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Analysis of windowing mechanisms with infinite-state stochastic Petri nets.

In this paper we present a performance evaluation of windowing mechanisms in world-wide web applications. Previously, such mechanisms have been studied by means of measurements only, however, given suitable tool support, we show that such evaluations can also be performed conveniently using infinite-state stochastic Petri nets. We briefly present this class of stochastic Petri nets as well as the approach for solving the underlying infinite-state Markov chain using matrix-geometric methods. We then present a model of the TCP slow-start congestion avoidance mechanism, subject to a (recently published) typical worldwide web workload. The model is parameterized using measurement data for a national connection and an overseas connection. Our study shows how the maximum congestion window size, the connection release timeout and the packet loss probability influence the expected number of buffered segments at the server, the connection setup rate and the connection time.

Introduction W HEN modeling and evaluating the performance of modern distributed systems, complex system behavior (involving networks, switches, servers, flow control mechanisms, etc.) as well as very complex workloads (often a mix of batch data, interactive data and real-time data for voice and video) need to be taken into account to come up with trustworthy performance measures. This most often leads researchers to either use a measurement-based approach or simulation as performance evaluation technique. As both these techniques are rather costly, it is worthwhile to study the suitability of analytic/numerical approaches based on stochastic Petri nets (SPNs) as well. As a challenging application for such a suitability study we have chosen the handling of WWW (World Wide Web) requests by the hypertext transfer protocol (HTTP) [1]. Clearly, there are many factors involved in this process: the client issuing the request, the server handling the re- quest, the type of request (just text, text with embedded graphics, pictures, or even video), the Internet connecting the client and the server, as well as the transport protocols used by client and server (TCP/IP). Taking all these aspects into account would render an analytical solution impossible, as also witnessed by the many measurement-based performance studies in this area ([2], [3]). To the best of the knowledge of the authors, no analytical performance studies for such systems have been reported yet. In this paper we propose to use a special class of SPNs for the construction and efficient numerical solution of per- The authors are with the Laboratory for Distributed Sys- tems, RWTH Aachen, 52056 Aachen, Germany. E-mail: farost,[email protected] formance models for the handling of WWW requests. Our models include client characteristics (request pattern) and request types, the Internet delays, the server speed, and the influence of the underlying TCP/IP protocol. In par- ticular, our model includes the explicit connection set-up (and release) phase of TCP/IP as well as its congestion avoidance windowing mechanism (known as slow start). Despite all these details, after a suitable abstraction pro- cess, the models can still be solved efficiently with numerical means, thereby using the tool spn2mgm to exploit the special quasi-birth-death (QBD) structure of the stochastic process underlying the SPN. The contribution of this paper is twofold. First, it shows the suitability of the SPN-based approach for the performance evaluation of complex systems, provided a suitable model class and solution method is chosen. Secondly, the performance results (obtained with realistic parameters derived from measurements) show the impact of lower-layer protocols (TCP/IP including its windowing mecha- nism) on the perceived performance at the application level (WWW). This paper is further organized as follows. In Section II, we concisely describe the class of SPNs and the employed solution methods, as well as the tool support we have de- veloped. We then describe the handling of WWW requests in detail, as well as the corresponding SPN model in Section III. Section IV is then devoted to a wide variety of numerical case studies. Section V concludes the paper. II. The Modeling Framework The basis of our modeling framework is a special class of SPNs (so-called infinite-state SPNs), which is suitable for an efficient numerical analysis. We present the main properties of these SPNs in the following section. Then, we discuss some issues concerning the numerical solution of the underlying Markov chain, and finally present the tool we developed to support infinite-state-SPN based modeling and analysis. A. Matrix-Geometric Stochastic Petri Nets Infinite-state SPNs are especially suited for modeling systems which involve large buffers or queues. Based on the class of (generalized) SPNs proposed by Ciardo in [4], they allow one distinguished place of unbounded capacity (graphically represented as a double-circled place) and represent an extension of the initial work by Florin and Natkin [5]. The unbounded place is usually used to model infinite buffers, or to approximate large finite buffers. While all features of Ciardo's original SPN class are still __ Fig. 1. Leveled structure of the underlying Markov chain of an infinite-state SPN. available (like immediate transitions, enabling functions, inhibitor arcs and arc multiplicities), there are some restrictions concerning the unbounded place: ffl The multiplicity of incoming and outgoing arcs is limited to one. ffl Transition rates and weights of immediate transitions may not depend on the number of tokens in the unbounded place. Due to the unbounded place, the Markov chain underlying the Petri net has an infinite number of states. The advantage of using infinite-state SPNs instead of classical ones is that the infinite state space has a special structure, which makes it eligible to very efficient numerical solution techniques, leading much quicker to a steady-state solution than by investigating a large, finite state space. B. Numerical Solution The continuous-time Markov chain underlying an infinite-state SPN is a QBD process [6], [7]. Its state space can be represented two-dimensionally as a strip which is unbounded in one direction. In the case of infinite-state SPNs, the position of a state in the unbounded direction corresponds to the number of tokens in the unbounded place. All states which belong to markings having the same number of tokens in the unbounded place are said to be in one level. Due to the restrictions mentioned in the previous section, transitions can only take place between states which belong to the same level, or between states in adjacent levels (see Fig. 1). Furthermore, the transition rates between levels are independent of the level index (except for a boundary level). This leads to an (infinite) generator matrix with the following block-tridiagonal structure: Except for the boundary, all transition rates can be kept in three square matrices A . The size of these matrices equals the number of states in the (non-boundary) levels (denoted by N in Fig. 1). There exist several efficient techniques for deriving the steady-state solution with of these Markov chains. The boundary conditions involving B represent a normal set of linear equations, and the non-boundary part Fig. 2. Snapshot of an editing session with agnes. is reduced to the quadratic matrix polynomial 0: The solution of this polynomial for R is the core of the solution techniques. While most methods provide an iterative solution to this problem (like those presented in [6], [8], [9]), the approach suggested in [10] transforms the problem to an Eigenvalue problem. Though the latter method has some interesting properties, the iterative approaches proved to be numerically more stable for large models so far (see [11] for a comparison). The results presented in this paper were obtained using the LR method [8]. C. Tool Support We developed extensive tool support to simplify the use of infinite-state SPNs for modeling and analysis, and for abstracting details of the underlying Markov chain. Using spn2mgm ([12], [13]), it is possible to specify the desired performance measures easily at the Petri net level in a reward-based way. Concerning the numerical solution, the user can choose between the algorithms mentioned in the previous section; all of them have been implemented using high-performance linear algebra packages. For the easy graphical specification of the Petri net, we adopted the generic net editing system agnes [14] (see Fig. 2). In addition, also a textual specification is still possible III. Modeling Windowed Traffic Control Mechanisms The TCP protocol relies on two window traffic control mechanisms to handle flow and congestion control. We focus on TCP's congestion control mechanism, which is described in the following section. Then, an appropriate SPN model is proposed, after which we suggest several approaches for modeling the traffic which has to be delivered by TCP. Finally, we describe how the model parameters used in our experiments are derived. with slow start and window flow control with slow start and window flow control Web browser/ client system request for frame/image frame/image data response segments frame/image data Lower protocol layers Web server segment proc. overhead client request Fig. 3. Accessing WWW documents via TCP/IP. A. System Description A TCP connection is associated with two windows: a receiver window to synchronize the sender's speed with the speed a receiver is able to process incoming data, and a congestion window to avoid network overload. The amount of data sent is given by the minimum of both windows. Assuming that the receiving station can handle incoming traffic sufficiently fast, the size of the congestion window is the limiting factor. avoids network congestion by limiting the number of packets traveling through the network simultaneously. This is accomplished by introducing the congestion win- dow, holding the remaining tolerable amount of data which can be fed into the network. Once a packet is acknowl- edged, the available congestion window space is increased by the appropriate number of bytes. Since the tolerable number of packets which can be fed into the network prior to being acknowledged is not known a priori, TCP adopts the slow start mechanism suggested by Jacobson [15]. Slow start initializes the congestion window size to one packet. Whenever a packet gets acknowledged, the window size is increased by one packet. Thus, the window size effectively grows exponentially. This exponential growth phase is bounded by a maximum window size parameter. Once it is exceeded, the congestion window size increases in a linear fashion. If a packet gets lost, the maximum window size parameter is set to half the current window size, and the current window size is set to one packet. While slow start provides a suitable algorithm for congestion avoidance, the initial phase of "probing for band- width" becomes a problem if a TCP connection is set up for transferring only a small amount of data. In particu- lar, the HTTP protocol suffers from this fact, since getting a HTML document and the images referenced therein is accomplished by building separate TCP connections for each of them (see Fig. 3). Since the amount of data per item is typically small, the slow start mechanism leads to a situation where just a fraction of the available band-width is used. Furthermore, the establishment of each connection involves the usual TCP three-way handshake, increasing delay by one round trip time. In version 1.1 of arr buffer no_conn conn server connect timeout tokens cwin ack ack_2 #cwin < max_win reset net loss loss_rate lost loss_done22 or rate *= #net rate *= #net rate *= #net Fig. 4. Petri net for modeling TCP behavior. the HTTP protocol, the use of persistent connections for several requests (P-HTTP) is possible and recommended, thus minimizing the number of necessary slow start phases and connection setups. Other approaches like T/TCP [16] and UDP-based mechanisms like ARDP [17] have been addressed in [3]. B. Model Development The model presented in this section has been developed having three main aims in mind. First, the model should account for the window flow control mechanism and capture the influence of the slow start mechanism on transmission performance. Second, connection setups have to be considered, since we are interested in the gain obtained by using persistent protocol versions. Third, the complexity of the overall model should not exceed the numerical capabilities of our analysis tool, so less important details had to be omitted. The model concentrates on the server-side of a client-server relationship. We assume that the main amount of traffic arises from server to client (as in the HTTP case), thus we neglect the impact of flow- and congestion control mechanisms on traffic which is sent to the server. Instead, different types of clients will lead to different patterns of generating segments (or packets) to be processed and transferred to the client. For the time being, we assume that the generation takes place according to a Poisson process. The model is illustrated in Fig. 4. The left hand side of the model deals with connection setup, while the right hand side represents windowing and the slow start mechanism. The segments to be delivered to the client are generated by the transition arr and put in the unbounded buffer place buffer. If there is no connection between server and client yet (as indicated by a token in no conn), transition server is disabled due to its enabling function and a connection setup is performed by firing transition connect. The connection will not be released until all segments in buffer have been delivered (inhibitor arc to timeout). Once the buffer is empty and all segments have been transferred (place net empty), the connection will be released after an additional delay, modeled by transition timeout. Each segment in buffer is the result of some operation of the server. In the HTTP case, the server has at least to perform some kind of database (or disk) access to retrieve the desired document. This overhead is modeled by the transition server. connection-setup and server overhead have been accounted for now, segments are ready for transmission. A prerequisite for submitting a segment to the network is that the congestion window is large enough. The total size of the congestion window is reflected by the place cwin. It initially holds one token, representing the size of one segment. The available size of the congestion window which currently remains for transmission is represented by the place tokens. Each segment submitted to the network takes one token from this place. After the segment has been acknowledged successfully (represented by transitions ack and ack 2), the token is put back. Transition ack 2 is enabled if the maximum allowed window size (max win) has not been reached yet. In this case, two tokens are put back into place tokens, and one additional token is put into cwin to reflect the enlarged congestion window. Transition ack is only enabled if the maximum congestion window size has been reached. In this case, the congestion window is not enlarged any further, and just one token is put back into tokens (we do not account for a further linear increase as featured by the original slow start procedure). As a last possibility, a segment may be lost in the network, represented by transition loss. In that case, the segment is put back in the buffer (this implies that the corresponding lost segment experiences the server delay induced by server for a second time; for excessive packet loss ratios, the server rate has to be adjusted appropriately) and the congestion window size is reset to one by enabling the immediate transition reset (actually, the congestion window is only reduced by the size of tokens still available in the congestion window to limit the size of the model). This transition is also enabled if no connection exists, thus effectively initializing the congestion window for the next connection. The exponentially distributed timing of transitions in the model surely is an approximation of the real-world system; it may be appropriate to model network delays, but time-out values usually have a less stochastic character. The impact of correctly modelling deterministic timeout values is, however, out of the scope of this paper; see e.g. [18] for an alternative to model connection release timeout values by Erlangian distributions. Note that the firing rates of ack, ack 2 and loss are proportional to the number of segments submitted to the network (place net). This approximation of an infinite- server behavior for network delays accounts for the fact that the increase of network load by submitting single segments is negligible. In future models, network delays could also be described more accurately, e.g. by using appropriate phase-type distributions as suggested in [19]. The most striking difference between this model and the 1 Note that this represents the overhead per segment of the server reply, not the overhead per client request. Therefore, per-request overhead has to be split between the number of segments resulting from this request. idle start frame_s img_s frame_m img_m select_s f1 leave_idle select_m weight_m im_rate buffer weight_s Fig. 5. Traffic generation model for HTTP. real slow start mechanism is that the maximum congestion window is set to a constant size (occurring as max win in the model), instead of setting it to half its value each time a segment is lost. Introducing an additional place for holding the current maximum window size would have led to a too large state space and has thus been omitted. Since the maximum window size can not shrink, the results obtained by analyzing the presented model provide an optimistic approximation. C. Workload Models As mentioned in the previous section, the slow start mechanism leads to poor bandwidth utilization when there is only a small amount of data to be transferred per TCP connection. In order to evaluate the merits of persistent connection approaches (like P-HTTP) in this case, an appropriate traffic model has to be developed. We focussed on modeling HTTP workload, using a characterization similar to the one presented in [3]. A user request for a web page is satisfied in two successive steps. First, the HTML document referenced by the URL is fetched from the server. Afterwards, all images referenced in this "frame" document are requested from the server. The traffic model must appropriately generate the segments which correspond to the server replies to these individual requests, and put them in the place buffer (Fig. 4). The traffic model we propose is shown in Fig. 5. It is able to account for two different user request types (a small and a medium one, see next section for details). After an exponentially distributed user-idle time has passed (tran- sition leave idle), a probabilistic choice between small and medium request type takes place. Each request type consists of two phases, corresponding to generating segments to satisfy the frame request and segments belonging to the image requests. While the duration of each phase corresponds to the time needed to submit the corresponding request to the server, the rate at which segments are generated during that phase (transitions T1,T2,T3 and T4) corresponds to the number of segments the server will send in reply to the request. buffer idle frame_s img_s leave_idle wait T3 Fig. 6. Arrival model for one request type with blocking after submitting a frame request. As an alternative to the arrival model shown in Fig. 5, we also investigate simplified versions of it, consisting of just one request type, simple IPP arrivals, and Poisson ar- rivals. Since these models do not account for the fact that prior to submitting an image request the reply to the preceding frame request has to be completed, we also investigate a blocking arrival model (see Fig.6). It deals with one request type only, where all frame segments have to be delivered before image segments are generated. This is realized by introducing an additional immediate transition which is disabled as long as there are any segments left in the server's outgoing buffer. D. Parameterization The parameters for the overall model can be split in three groups: server, network, and workload parameters (see Table I). The following paragraphs explain how the SPN transition rates can be derived from these. D.1 Server Performance Server-specific performance data is introduced in the model by the transition server. Its mean firing time corresponds to the average workload per segment of an answer to a client request. We assume that it is due to three param- eters: computational effort per request, disk seek time per request and disk transfer time per segment. Per-segment workloads are obtained by dividing the per-request overheads by the mean number of segments per request, denoted by s mean (derived in the workload parameterization section). In conclusion, the firing rate of transition server is given by s mean disk D.2 Network parameters Apart from the usual packet transfer latencies, a connection setup requires an additional round trip time to account for the TCP three-way handshake, so the rate of connect is rtt . The rate of the connection release transition timeout is given by t \Gamma1 release . The time needed to acknowledge a segment sent from server to client consists of one round trip time plus the segment transfer time, which depends on segment size Server characteristics Computation time per request [s] t comp Disk transfer speed [B/s] disk Network characteristics Round Bandwidth [B/s] bw Loss ratio p loss Max. TCP segment size [B] nMSS Connection release timeout [s] t release Max. congestion window size [segs] max cwin Workload characteristics User idle time [s] t idle Small request probability p small Small frame request size [B] b sf Number of small image requests n si Small image request sizes [B] b Medium frame request size [B] b mf Number of medium image requests n mi Medium image request sizes [B] b I Overview on all model parameters. and bandwidth. Since we also consider packet losses, the rates of transitions ack and ack 2 are given by loss ), and the rate of transition loss equals loss . D.3 Workload parameters Due to the small size of client requests, we assume that the submission of a request takes on average one round trip time. Thus, the firing rates of transitions f1 and f2 equal rtt . We also assume that all image requests are submitted simultaneously, thus again involving just one round trip time and yielding the rate t \Gamma1 rtt for transitions i1 and i2 as well. The number of segments to be put into the server's buffer place depends on the size of the reply corresponding to a request. For small requests, it is given by s db sf =nMSS e for the initial frame request. The corresponding total number of segments to be transferred in reply to the image requests is s db si;k =nMSS e. These are the numbers of segments to be generated on average while a token is in places frame s and img s, so the rates of and are s sf =t rtt and s si =t rtt , respectively. The parameters for the medium size request type can be computed similarly. Also, using n si , n mi and p small , the mean number of segments per request can be computed as Parameter National International loss nMSS 536 [B] II Network parameters used in the experiments. IV. Performance Evaluation The SPN model proposed in the previous section features a wealth of model parameters, and a large number of performance measures can be obtained from its analysis. We will thus keep many parameters constant throughout all experiments, and focus on the investigation of a few aspects only. We present the numerical results after describing the selected parameters in the next section. Fi- nally, some information on the model's complexity and the computational solution effort are given. A. Parameter selection Since the following investigations focus on the variation of a few parameters only, many of the model parameters given in Table I are kept constant. Server characteristics. We assume that the computation time per request is t 0:01[s]. The performance parameters of the disk are t Network characteristics. Parameters concerning network performance have been taken from [19], where extensive Internet performance investigations have been accomplished. We selected two reference connections, a national one (RWTH Aachen to University of Karlsruhe, Germany) and an international connection (RWTH Aachen to Stanford University, U.S.A. See Table II for the parameters of these connections. If not mentioned otherwise, the connection release time-out t release has been set to 10 seconds. Workload characteristics. The parameters for small and medium request types occurring in the model's workload characterization have been taken from [3], representing the structure of some popular Web pages. The small request type consists of an initial 6651 byte frame page, referencing two images of size 3883 and 1866 bytes. Medium requests are formed by a 3220 byte frame and three images of size 57613, 2344 and 14190 bytes. We assume that the probability for a small request is B. Numerical Results Our investigations focus on three main areas. We first investigate the impact of different workload models on the obtained performance measures. Here, we were also interested to which extent the maximum congestion window size (max cwin) influences the performance characteristics.501502503502 4 mean buffer size maximum congestion window size full HTTP workload model IPP workload model one request type only one request type/blocking model Poisson model Fig. 7. Expected buffer size for the international connection and different workload models.5152535 mean buffer size maximum congestion window size full HTTP workload model IPP workload model one request type only one request type/blocking model Poisson model Fig. 8. Expected buffer size for the national connection and different workload models. The connection release timeout plays an important role when dealing with protocols like persistent HTTP. We thus show how changing this parameter affects the properties of the overall system in the next experiment. Finally, the influence of the segment loss probability on the system has been investigated. B.1 Workload models and congestion window size As a first point, we were interested in how far detailed workload models influence the results of our investigation. Since the workload model greatly enlarges the number of states of the Markov chain underlying the SPN (e.g., the QBD level size of the arrival model shown in Fig. 5 is five times as large as a model with simple Poisson arrivals), it is interesting to see whether this effort pays off. Fig. 7 shows the mean buffer size (number of tokens in place buffer) for different workload models. Clearly, increasing the maximum window size leads to a higher segment throughput, and thus reduces the buffer filling. connection probability maximum congestion window size Poisson model full HTTP workload model one request type only IPP workload model one request type/blocking model Fig. 9. Probability for an existing connection for different workload models. Concerning the different workload models, the results for the full HTTP model (as shown in Fig. 5) differ significantly from those of the simplified versions. The results of the one-request and IPP workload models (with identical mean segment arrival rate) are very similar and still capture the qualitative behavior of the original model. How- ever, due to ignoring the bursty arrival pattern accounted for by the other workload models, the approximation of the workload by a Poisson process leads to a dramatic under-estimation of the expected buffer size. As an interesting point, the results for the blocking workload model (Fig. do not differ too much from the non-blocking one-request- type model, especially for larger maximum connection window sizes. The absolute values are smaller, since the mean segment generation rate for this workload model is lower than for the other models (due to the additional waiting time in place wait). Fig. 8 illustrates the same performance measures for the national Internet connection. Due to the lower round trip time, the average buffer filling is much lower than in the international case. Again, the Poisson workload model yields much too low results. From now on focusing on the international connection type in all experiments, we investigate the steady-state probability for an existing connection (i.e., a token in place conn) in Fig. 9 . While the results for all bursty workload models coincide, the Poisson workload distributes segments much more in time, leading to a situation where the connection timeout almost never expires. Due to the increased usable bandwidth for larger maximum connection windows, the segments in the server's buffer are delivered quicker to the client, which leads to the connection to be released quicker. This results in lower connection probabilities for larger values of max cwin. Fig. 10 shows the connection setup rate for the different arrival models, which may represent a cost-factor. As can be observed, the setup rate increases for larger congestion window sizes, since requests are satisfied quicker and the0.010.020.030.04 connection setup rate maximum congestion window size one request type/blocking model full HTTP workload model one request type only IPP workload model Poisson model Fig. 10. Connection setup rates for different workload models.0.30.50.70.90 connection probability connection release timeout value [s] Poisson model one request type only one request type/blocking model Fig. 11. Probability for an existing connection for different connection release timeouts. release timeout expires more often in this case. Again, the Poisson model leads to significantly different results. Summarizing, it can be said that ignoring the work-load burstiness dramatically alters the performance measures obtained from the model, however, thanks to our modeling environment, bursty arrival patterns can easily be accounted for. Furthermore, increasing the maximum congestion window above a minimum value of about 8- segments leads to much better bandwidth utilization, thus reducing the buffer size and the time connections are held, albeit at the cost of increased connection setup rates. Though increasing the bandwidth utilization by higher maximum congestion window sizes is generally desir- able, this may also increase network congestion and packet losses. However, this aspect can not be considered with the single client-server model presented here. 0.10.20.30.4 connection setup rate connection release timeout value [s] Poisson model one request type only one request type/blocking model Fig. 12. Connection setup rates for different connection release time-outs B.2 Influence of connection release timeout. The introduction of a connection release timeout is crucial for the reduction of connection-setups and delays when protocols like P-HTTP are employed. Clearly, when choosing this parameter it is important to compare the gain of less connection setups with the higher costs imposed by maintaining a (mainly unused) connection. Fig. 11 illustrates the probability of an existing connection for different values of t release for a maximum connection window size of 12. Obviously, the probability increases for larger timeout values. Since the amount of data to be transferred remains constant, an existing connection is often unused. On the other hand, the connection setup rate decreases for larger timeout values, as illustrated in Fig. 12. B.3 Influence of segment loss ratio Apart from bandwidth and average round trip time, the packet loss probability heavily influences the performance of an Internet connection. The impact of losses on the windowing system is shown in Fig. 13 for different values of max cwin. It can be observed that the mean buffer size of the system increases dramatically for high loss ratios if cwin is chosen too small. The system's behavior is much more robust concerning packet losses if the connection window size is large enough. This behavior is due to the fact that higher packet loss ratios effectively increase the amount of segments to be delivered by the network, since lost segments are re-submitted for transmission by transition loss done. For a loss ratio of 0:5, every second packet has to be retransmitted. Since lost packets are again subject to loss in the next transmission try, the number of segments to be delivered effectively doubles. Since small values of max cwin lead to small effective transmission performance, the buffer size is particularly sensitive to packet losses.20060010000 0.1 0.2 0.3 0.4 0.5 mean buffer size segment loss probability full HTTP workload model, max_cwin=4 full HTTP workload model, max_cwin=8 full HTTP workload model, max_cwin=12 Fig. 13. Expected buffer size for different loss probabilities. C. Computational effort The Markov chain underlying the investigated SPN can grow remarkably large. For example, the SPN shown in Fig. 4 with the full workload model as shown in Fig. 5 leads to an underlying QBD process with 765 states per level. In some experiments (see e.g. Fig. 13), we obtained mean buffer sizes around 1000. Consequently, the evaluation of this system by using a large finite Markov chain would involve the investigation of several thousand levels, leading to a total number of several million states. While deriving steady-state measures of Markov chains of this size becomes a problem when using common numerical or simulation methods, the QBD-based solution approach leads to results in a quick and memory-efficient way. We were able to accomplish the solution of the above-mentioned model in around 2.5 hours on a SUN SparcStation 20 clocked at 75 MHz. This time includes generating the state space, recognizing the QBD structure, solving the QBD Markov chain, and computing the desired performance measures. However, it should be noted that models of this size currently represent the upper limit we are able to solve due to numerical instabilities. V. Conclusions To our best knowledge, we presented the first performance evaluation study of the TCP slow-start mechanism under P-HTTP workload which is based on the numerical analysis of a SPN model. It has been shown that the choice of a reasonably high limit for the maximum congestion window is crucial for efficiently utilizing the communication infrastructure. This is especially true for connections with high packet losses. We also demonstrated the strong influence of the employed workload model on the results of the system anal- ysis. In particular, the approximation of bursty workloads by a Poisson process yields misleading results. Our modeling environment, based on powerful yet user-friendly tools for using efficient numerical techniques and hiding them behind a SPN-based interface proved to be of great use in the experiments. By employing QBD-based methods for the analysis of the underlying Markov chain we were able to derive numerical results much quicker than by conventional methods. Concerning the application investigated here, far more extensive experiments are possible. We did not present the impact of varying the speed of the server or the impact of changing workloads and user behavior. Further investigations could also focus on derivations of the slow-start algorithm. Of course, while the results presented here are reasonable, the validation of (especially more detailed) models by measurements and simulations is an important topic. Future work will also focus on the improvement of our modeling environment. In particular, the limit of about 1000 states in the repeating levels of the underlying QBD process is often a problem (for example, it led to the relatively small choice of the maximum congestion window parameter in the presented experiments). Future work will concentrate on the combination of sparse matrix computations with the spectral expansion solution method for QBD processes [10] to alleviate this problem. Furthermore, this method would also allow us to drop the first requirement on the infinite SPN place mentioned in Section II. Acknowledgments This work has been supported by the doctorate program computer science and technology at the RWTH Aachen. We also thank the department of Prof. Hromkovic for donating some spare computing resources. --R "Hypertext transfer protocol - HTTP/1.1," "Internet Web Servers: Workload Characterization and Performance Implications," "Modeling the performance of HTTP over several transport protocols," "SPNP: Stochastic Petri net package," "One place unbounded stochastic Petri nets: Ergodicity criteria and steady-state solution," Matrix Geometric Solutions in Stochastic Mod- els: An Algorithmic Approach "Matrix geometric solutions in Markov models: A mathematical tutorial," "A logarithmic reduction algorithm for quasi birth and death processes," "Analysis of a finite capacity multi-server delay-loss system with a general Markovian arrival process," "Spectral expansion solution for class of Markov models: Application and comparison with the matrix-geometric method," "Steady-state analysis of infinite stochastic Petri nets: A comparison between the spectral expansion and the matrix-geometric method," "SPN2MGM: Tool support for matrix-geometric stochastic Petri nets," spn2mgm Web Page "Entwurf und Implementierung einer parametrisierbaren Benutzeroberflache fur hierarchische Netz- modelle," "Congestion avoidance and control," "Extending TCP for transactions - concepts," The virtual system model: A scalable approach to organizing large systems "Untersuchungen zum Verbindungsmanagement bei Videoverkehr mit Matrix- geometrischen stochastischen Petrinetzen," Messung und Modellierung der Dienstgute paketvermittelnder Netze --TR

matrix-geometric methods;window flow control;congestion control;stochastic Petri nets

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Filter-based model checking of partial systems.

Recent years have seen dramatic growth in the application of model checking techniques to the validation and verification of correctness properties of hardware, and more recently software, systems. Most of this work has been aimed at reasoning about properties of complete systems. This paper describes an automatable approach for building finite-state models of partially defined software systems that are amenable to model checking using existing tools. It enables the application of existing model checking tools to system components taking into account assumptions about the behavior of the environment in which the components will execute. We illustrate the application of the approach by validating and verifying properties of a reusable parameterized programming framework.

INTRODUCTION Modern software is, increasingly, built as a collection of independently produced components which are assembled to achieve a system's requirements. A typical software system consists of instantiations of generic, reusable components and components built specifically This work was supported in part by NSF and DARPA under grants CCR-9633388, CCR-9703094, and CCR-9708184 and by NASA under grant NAG-02-1209. for that system. This software development approach offers many potential advantages, but it also significantly complicates the process of verifying and validating the correctness of the resulting software systems. Developers who wish to validate or verify correctness properties of software components face a number of challenges. By definition, reusable components are built before the systems that incorporate them, thus detailed knowledge of the context in which a component will be used is unavailable. Components are oftened designed to be very general, for breadth of applicability, yet configurable to the needs of specific systems; this generality may impede component verification. Typ- ically, components are subjected to unit-level testing and are delivered with informal documentation of the intended component interface behavior and required component parameter behavior. For high-assurance systems this is lacking in a number of respects: (i) unit-level reasoning focuses solely on local properties of the component under consideration, (ii) informal documentation cannot be directly incorporated into rigorous reasoning processes, and (iii) system developers may have some knowledge about the context of component use, but no means of exploiting this information. In this pa- per, we describe an automatable approach to applying existing model checking tools to the verification of partial software systems (i.e., systems with some missing components) that addresses these concerns. Model checking is performed on a finite-state model of system behavior not on the actual system artifacts (e.g., design, code), thus any application of model checking to software must describe model construction. In our approach, models for partial systems are constructed in two independent steps. First, a partial system is completed with a source code representation of the behavior of missing system components. This converts the open partial-system to a closed system to which model checking tools can be applied. Second, techniques from partial evaluation and abstract interpretation [19, 18] are applied to transform the completed source code into the input format of existing finite-state system generation tools. Finite-state models built in this way are safe, thereby insuring the correctness of verification results, but may be overly pessimistic with respect to the missing components' behavior. To enhance the precision of reasoning, we filter [16] missing components' behavior based on assumptions about allowable behavior of those components. Our approach supports model checking of systems with different kinds of missing components, including components that call, are called from, and execute in parallel with components of the partial sys- tem. This flexibility enables verification of properties of individual components, collections of components, and systems. The work described in this paper extends the applicability of existing model checking techniques and tools to partial software systems and illustrates the practical benefits of the filter-based approach to automated anal- ysis. We illustrate our approach and its benefits by verifying correctness properties, expressed in linear temporal logic (LTL) [24], of realistic generic, reusable com- ponents, written in Ada, using the SPIN model checker [20]. In principle, the approach described in this paper can be used in any setting that supports filter-based analysis [16]. In the following section we discuss relevant background material. Section 4 discusses abstractions used in constructing finite-state software models and Section 5 describes our approach to completing partial software sys- tems. We then present our experiences and results from applying the analysis approach to a generic, reusable software component in Section 6. Section 7 describes related work and we conclude, in Section 8, with a summary and plans for future work. In this section we overview LTL model checking, a variant of model checking for open systems called module checking, and a technique for refining model checking results using filter formulae. These ideas form the basis for our approach to constructing and checking finite-state models of partial software systems. 3.1 Model Checking Model checking techniques [7, 20] have found success in automating the validation and verification of properties of finite-state systems. They have been particularly effective in the analysis of hardware systems [26] and communication protocols [20, 30]. Recent work has seen model checking applied to more general kinds of software artifacts including requirements specifications [2, 3], architectures [28], and implementations [13, 9]. In model checking software, one describes the software as a finite-state transition system, specifies system properties with a temporal logic formula, and checks, exhaus- tively, that the sequences of transition system states satisfy the formula. There are a variety of temporal logics that might be used for coding specifications. We use linear temporal logic in our work because it supports filter-based analysis and it is supported by robust, efficient model checking tools such as SPIN [20]. In LTL a pattern of states is defined that characterizes all possible behaviors of the finite-state system. We describe LTL operators using SPIN's ASCII notation. LTL is a propositional logic with the standard connectives &&, -?, and !. It includes three temporal operators: !?p says p holds at some point in the future, []p says p holds at all points in the future, and the binary pUq operator says that p holds at all points up to the first point where q holds. An example LTL specification for the response property "all requests for a resource are followed by granting of the resource" is [](request -? !?granted). SPIN accepts design specifications written in the Promela language and it accepts correctness properties written in LTL. User's specify a collection of interacting processes whose product defines the finite-state model of system behavior. SPIN performs an efficient non-empty language intersection test to determine if any state sequences in the model conform to the negation of the property specification. If there are no such sequences then the property holds, otherwise the sequences are presented to the user as exhibits of erroneous system behavior. 3.2 Module Checking In computer system design, a closed system is a system whose behavior is completely determined by the state of the system. An open system (or module [22, 23]) is a system that interacts with its environment and whose behavior depends on this interaction. Given an open system and temporal logic formula, the module checking problem asks whether for all possible environments, the composition of the model with the environment satisfies the formula. Fortunately, for LTL, the module checking problem coincides with the basic model checking problem[22]. Often, we don't want to check a formula with respect to all environments, but only with respect to those that satisfy some assumptions. In the assume-guarantee paradigm [29], the specification of a module consists of two parts. One part describes the guaranteed behavior of the module; which we encode in the finite state system to be analyzed. The other part specifies the assumed behavior of the environment with which the module is interacting and is combined with the property formula to be analyzed. 3.3 Filter-based Analysis Filters [16] are constraints used to incrementally refine a naively generated state space and help validate properties of the space via model checking. Filters can be represented in a variety of forms (e.g., as automata or temporal logic formulae) and are used in the FLAVERS static analysis system [14]. Filters in FLAVERS were originally developed to sharpen the precision of analysis relative to internal components of a complete software system that were purposefully modeled in a safe, but abstract manner. Filters serve equally well in refining analysis results with respect to external component behavior in the analysis of partial software systems [11]. In this paper, we encode filters in LTL formulae to perform assume-guarantee model checking. Given a property that encode assumptions about the environment, we model check the combined We refer to the individual F i as filters and to the combined formula as a filter formula. We note that many specification formalisms support filter-based analysis, but some popular formalisms, such as CTL, do not. It may be possible to encode a filter into a CTL formula. In general, however, because multiple temporal operators cannot lie directly in the scope of a single path quantifier, there is no simple method for constructing a filter formula in CTL. In principle, model checking can be applied to any finite-state system. For non-trivial software systems we cannot render a finite-state system that precisely models the system's behavior, since, in general, the system will not be finite-state. Even for finite-state software, the size a precise finite-state model will, in general, be exponential in the number of independent components, i.e., variables and threads of control. For these rea- sons, we would like to use finite-state system models that reflect the execution behavior of the software as precisely as possible while enabling tractable analysis. We use techniques from abstract interpretation to construct such models. In the remainder of this section, we introduce the notion of abstract interpretation, we then describe a collection of abstract interpretations that we use in model construction, and finally, we describe our approach to the selection of abstractions to be used for a given software system. 4.1 Safe Models for Verification We say that a finite-state model of a software system is safe with respect to model checking of a property specification if model checking succeeds only when the property holds on the real system. For LTL model- checking, which is fundamentally an all-paths analy- sis, any abstraction of behavior must preserve information about all possible system executions. This class of abstractions can be described as abstract interpretations (AI) [10] over the system's execution semantics. These abstractions are similar to the kinds of approximations that are introduced into program representations (e.g., control flow graphs) used in compiler analyses [27]. When behaviors are abstracted in this way and then exhaustively compared to an LTL specification and found to be in conformance, one can be sure that the true executable system behaviors conform to the specification. One of the strengths of AIs is that they guarantee the safety of information gathered by analyses that incorporate them. To achieve this we need to precisely define each AI. We formalize an AI as a \Sigma-algebra [18] which defines, for a concrete type signature in the source pro- gram, a type for the abstract domain of the AI and abstract definitions for the operations in the signature. Operationally we view a \Sigma-algebra as a data abstraction with a defined domain of values and implementations of operations over that domain. This operational view allows systematic abstraction by substituting abstract definitions for concrete definitions for abstracted program variables. Computation with abstract values and operations can then proceed in the same way as it would have for concrete values and operation. Unlike the concrete operations in a program, we can define abstract operations to produce sets of values of the operation's return type. This is a mechanism for encoding lack of precise information about variable values. This mechanism can be exploited by the partial evaluation capabilities discussed in Section 5. The partial evaluator will treat a set of values returned by an operation as equally likely possibilities and create a variant program fragment for each value (i.e., simulating non-deterministic choice). This allows subsequent analyses (e.g., model checking) to detect the presence of specific values that may be returned by the operation. The technical details of how this is performed is given in [18] for a simple imperative language; we have scaled those techniques up and applied them to a real Ada program in Section 6. Conceptually, one can think of partial evaluation as the engine that drives the systematic application of selected abstract interpretations to a given source program. 4.2 Sample Abstract Interpretations Space constraints make it impossible to give the complete formalization of the AIs used in the example in Section 6. In this section, we describe the main idea of each AI and illustrate selected abstract operations. The point AI is the most extreme form of abstraction. Under this AI, a variable's abstract domain has a single value, representing the lack of any knowledge about possible variable values. Abstract operations for the variable are defined as the constant function producing the domain value. Abstract relational operations that test the value of a variable are defined as the constant function returning the set ftrue; falseg. While the point AI is extreme in its abstraction, it is never- theless, not uncommon in existing FSV approaches- many state reachability analyses (e.g., CATS [32]) use this abstraction for all program variables. Closely related to the point AI is the choice AI. This AI also encodes a complete lack of knowledge about possible program variables, but it does so in a different way. Abstract operations are defined to produce the set of all possible domain values. Abstract relational operations that test the value of a variable retain their concrete semantics. In most cases, the possible values reaching such a test will consist of the set of all domain values and the result of the test will be the set ftrue; falseg. Exposing the distinct domain values to a partial evaluator gives it the opportunity to specialize program fragments with respect to possible variable values. For ex- ample, a control flow branch for each variable value can be introduced into the program, subsequent program fragments can be hoisted into each branch, and each fragment can be specialized to the branch variable value. Any tests of an abstracted variable within such a branch will have a single domain value flowing into it, thus, a more refined test result (e.g., true or false, but not both) may be computed. The resulting program model can be more precise than using the point AI, but care should be taken in applying the choice AI since it will result in a larger program model. The k-ordered data AI provides the ability to distinguish the identity of k data elements, but completely abstracts the values of those elements. For 2-ordered data AI, any pair of concrete data values are mapped to abstract values d 1 and d 2 ; all other values are mapped to the ot value. Like in the choice AI, abstract operations are defined to return sets of values to model lack of knowledge about specific abstract values. For all operations, except assignment, the constant function returning fd 1 ; d 2 ; otg is used. For assignment the identity function is used. Relational test operations are slightly more subtle. Relational operations other than equality, and inequality, return falseg. Equality is defined as: ffalseg if x 6= y Inequality is defined analogously. A special case of the classic signs AI [1] is the zero-pos AI which is capable of differentiating between valuations of a variable that are positive and zero. The abstract domain ranges over three values: unknown; zero; positive. For this AI we find it convenient to introduce the unknown value, which represents the fact that the variable can have either zero or positive value. Abstract operations for assignment of constant zero, increment and decrement by a positive value, and a greater-than test with zero are defined as: positive Other operations are defined analogously. It is also possible to define an AI that is safe with respect to a restricted class of LTL properties. One such abstraction is used commonly in verifying message ordering requirements in communication protocols. Wolper [31] has shown that reasoning about pairwise ordering questions over a communication channel that accepts large domains of values can be achieved using a domain of size three 1 . This can be achieved when the data in the channel are not modified or tested by the program 2 . We support reasoning for this class of questions through the use of a 2-ordered list AI. This AI represents the behavior of a "list" of data items which are themselves abstracted by the 2-ordered data AI. Conceptually, the values of the abstracted list record whether a specific d i has been inserted into the list and not removed yet. If both of the non-ot 2-ordered data are in the list their ordering is also recorded. No attempt is made to represent the number of ot elements in the list or their relative position with respect to the values. The abstract list values are: some : zero or more ot values mixed with zero or more ot values mixed with zero or more ot values mixed with zero or more ot values, with d 1 in front of d 2 and d 1 mixed with zero or more ot values, with d 2 in front of d 1 Technically, this AI is not safe for LTL (e.g., it does not allow for lists with multiple instances of d 1 values), however, the abstraction is safe for all system executions for which the d i are inserted at most once into the list. Thus, the AI is safe for LTL formulae of the Requirements involving more than a pair of data items can be handled by a simple scaling of the approach described here. This condition can be enforced using an approach that is similar to the restriction of errors discussed in Section 5. where P is an arbitrary LTL formula and the call and return prefixes indicate the program actions of invoking and returning from an operation. This is a filter- formula, as described in Section 3, that restricts checking of P to paths that are consistent with the information preserved by the the 2-ordered list AI. We illustrate abstract list operations for Inserting elements at the tail and Removeing elements at the head; other operations are defined analogously. f(ot; L)g otherwise Since list operations may both produce values and up-date the list contents we define abstract operations over tuples. The first component is the return value of the indicates that no value is returned. The rest of the components define how the components of the AI should be updated based on the operation. In the next section we discuss the use of a 2-ordered abstraction in completing a partial system that enables verification of ordering properties of data-independent systems. 4.3 Abstraction Selection Given a collection of program variables and a collection of AIs we must select, for each variable, the AI which will define its semantics in the finite-state model. We do not believe that this process can be completely automated in all cases. Our experience applying AIs in model construction, however, has left us with a methodology and a set of heuristics for selecting abstractions. In this methodology, we bind variables to AIs, then use additional program information to refine the modeling of a variable by binding it to a more precise AI. Start with the point AI Initially all variables are modeled with the point AI. Use choice for variables with very small domains Variables with domains of size less than 10 that are used in conditional expressions are modeled with the choice AI. semantic features in the specification The property to be checked includes, in the form of propositions, different semantic features of the program (e.g., valuations of specific program variables). These features must be modeled precisely by an AI to have any hope of checking the property. Select controlling variables In addition to variables mentioned explicitly in the specification, we begin if not d then if y?0 and not d then Figure 1: Example for AI Selection consider variables on which they are control de- pendent. The conditional expressions for these controlling variables suggest semantic features that should be modeled by an AI. Select variables with broadest impact When confronted with multiple controlling variables to model, select the one that appears most often in a conditional. After the selection process is complete, we generate a finite-state model using the variable-AI bindings and check the property. Model checker output either proves the property or presents a counter-example whose analysis may lead to further refinement of the AIs used to model program variables. To illustrate this methodology, consider the program fragment in Figure 1 which has variables d, x and y. Assume we are interested in reasoning about the response property [](xISzero -? !?call-P). The key features that are mentioned explicitly in this specification are values of variable x and calls to procedure P. We must model x with more precision than the point AI provides in order to determine the states in which it has value zero. An effective AI for x must be able to distinguish zero values from non-zero values; we choose the zero-pos AI. At this point we could generate an abstracted model and check the property or consider additional refinements of the model; we choose the latter for illustrating our example. Using control dependence information we can determine the variables that appear in conditionals that determine whether statements related to x and P execute. In our example, there are two such variables d and y. We could refine the modeling of both of these variables, but, we prefer an incremental refinement to avoid unnecessary expansion of the model. In choosing between these variables, we see that d appears in both conditionals and we choose to model it since it may have a broader impact than modeling y. Since, d is a boolean variable and the conditional tests for falsity, we choose for it to retain its concrete semantics. At this point, we would generate an abstracted model and check the property. If a true result is obtained then we are sure that the property holds on the program, even though the finite-state system only models two variables with any precision. If a false result is obtained then we must examine the counter-example produced by the model checker. It may reveal a true defect in the program or it may reveal an infeasible path through the model. In the latter case, we identify the variables in the conditionals along the counter-example's path as candidates for more binding to more precise AIs. This methodology is not foolproof. It is based on a fixed collection of AIs and a given program variable may require an AI that is not in that collection. Our heuristics for choosing variables to refine may cause the generation of finite-state models that are overly precise, and whose analysis is more costly than is necessary. Nevertheless, this approach has worked well on a variety of examples and we will continue to improve it by incorporating additional AIs and mechanisms for identifying candidates for refinement. While not fully-automatable, this methodology could benefit from automated support in computing controlling variables and in the analysis of counter-examples. We are currently investigating how best to provide this kind of support. A partial system is a collection of procedures and tasks 3 . We complete a partial system by generating source-code that implements missing components, which we call contexts. These context components are combined with the given partial system. Stubs and drivers are defined to represent three different kinds of missing components; calling, called and parallel contexts. Calling contexts represent the possible behavior of the portions of an application that invoke the procedures of the partial system. Called contexts represent the possible behavior of application procedures that are invoked by the procedures and tasks of the partial system. Parallel contexts represent those portions of an application that execute in parallel with and engage in inter-task communication with the procedures and tasks of the partial system. For simplicity, in our discussion and the examples in this section, we phrase our system models and properties in terms of events (i.e., actions performed by the software). It is often convenient to use a mixture of event and state-based descriptions in models and properties and we do so in the example in Section 6. To begin construction of any system model one must have a definition of the events that are possible. For Ada programs, these events include: entry calls or accepts, calls or returns from procedures, designated statements being executed or variables achieving a specified value. We partition the events into those that are 3 The approach in this section can easily be extended to support packages and other program structuring mechanisms. procedure P() is begin task body T is begin accept accept procedure stub() is choice begin loop case choice is when 1 =? A; when 3 =? P(); when 4 =? null; otherwise =? exit; task body driver is begin task body T is choice begin accept - call missing; loop case choice is when 1 =? A; when 3 =? null; ? when 4 =? null; otherwise =? exit; return missing; accept task body driver is choice begin loop case choice is when 1 =? A; when 3 =? T.E; C; when 4 =? null; otherwise =? exit; Figure 2: Partial Ada System, Stubs and Drivers internal to the components being analyzed and external events, which may be executed by missing components. Based on this partitioning we construct a stub procedure that represents all possible sequences of external events and calls on public routines and entries in the partial system. Figure 2 illustrates, on the left side, a partial system consisting of procedure P and task T. The internal events are calls on the E entry of T and execution of C. Two external events are defined as A and B. The stub procedure and driver task are also given in the figure. Because there are no external entries there is no parallel context defined. Existing model checking tools require a single finite-state transition system as input. To generate such a system from a source program with procedures requires inlining, or some other form of procedure integration. We describe the construction of a source-level model for a completed partial system as a series of inlining operations. We assume that there are no recursive calls in the system's procedures, stubs and drivers. Given this assumption, Figure 3 gives the steps to assemble a completed system. Applying these steps to the example gives the code on the right side of Figure 2. The same stub procedure is used to model the behavior of all missing called components. To enable model checking based on assumptions about the behavior of Input: collection of procedures, tasks and a description of the external alphabet Output: a source level system without external references Steps: 1. Generate stub procedure that non-deterministically chooses between the actions in the external alphabet and calls to the procedures and entries of the program components. It must also be capable of choosing to do nothing or to return. 2. Inline calls to procedures made by stubs. 3. Stubs may now contain (inlined) calls to task entries. For each task that calls a missing component specialize the stub so that any calls to that tasks entries are replaced by an error indication. 4. Calls to missing components from tasks are replaced by the stub routine. Indications of the call to and return from the missing component are inserted before and after the stub body. 5. The driver and parallel contexts, if any, are formed by in-lining the stub body. Figure 3: System Completion Algorithm specific missing components we bracket the inlined stub with indicators of call and return events for the missing component (e.g., call missing and return missing). The goal of the system completion process is to produce legal Ada source code so that subsequent tools can process the system. This is somewhat at odds with the fundamental lack of knowledge about event ordering in missing components. We model this lack of knowledge with non-determinism by introducing a new variable, choice, that is tested in the stub conditionals. This variable will be abstracted with a point AI and subsequent model construction tools will represent such conditionals with non-deterministic choice. We must take care to insure that potential run-time errors are preserved in the completed system, since they contribute to the actual behavior of the software. For example, it is possible for system tasks to call stubs which in turn call system procedures containing entry calls on that task; this is a run-time error in Ada. To preserve this possible behavior we introduce an error event. In the example in Figure 2 the point at which such an event is introduced is marked by a ?. This allows a user to test for the possibility of the run-time error or to filter the allowable behavior of missing components to eliminate the error (i.e., using []!error as the property to be checked or as a filter). This conversion is done separately for each task and amounts to specialization of the stub body for the task by interpreting self-entry calls as the error event. Completing a partial program does not yield a finite-state system. The next step is to selectively abstract program variables and transform dynamic program behavior to a static form. 5.1 Automating Model Construction Our approach to automating the construction of safe finite-state models of software systems builds on recent work in abstraction-based program specialization [18]. Figure 4 illustrates the steps in converting Ada source to Promela which can be submitted to the SPIN model checker. First, the partial system is completed with a source-level model of its execution environment. We apply a source-to-source partial evaluation tool that transforms the program to a form that is more readily modeled as a finite-state system. Partial evaluation is a program transformation and specialization approach which exploits partial information about program data. Essentially it performs parts of a program's computation statically; the result is a simplified program that is specialized to statically available data values. A wide-variety of source transformations can be applied to aid in finite-state model construction including; procedure integration, bounded static variation, and migration of dynamically allocated data and tasks to compile-time [19]. A novel feature of the approach we use is its ability to incorporate AIs for selected variable [18]; we use the variable-AI bindings discussed in Section 4. After partial evaluation, a tool is applied to convert the resulting Ada to SEDL, an internal form used by the Inca [4] toolset 4 , which is then converted into Promela. The Promela is then submitted, along with an LTL specification, to SPIN which produces either an indication of a successful model check or a counter-example. Aside from selection of AIs, this approach is completely automatable. At present, the system completion and the partial evaluation tools are not fully-implemented and were not used in the experiments described in Section 6; all other tools depicted in Figure 4 were run without user intervention. Stubs, drivers and the ab- stracted, specialized Ada for those experiments were constructed by-hand using the algorithms that are being implemented in our partial evaluation tool. We are implementing an approach to stub and driver generation using ideas from work on synthesis of program skeletons from temporal logic specification [25]. With this approach we will be able to encode filters on environment behavior directly into stubs and drivers, thereby eliminating the need for including those filters in the formula to be checked. It remains to be seen whether encoding filters in the transition system or in the formula to be checked results in better performance; we plan to explore this question in future work. 4 Inca was previously referred to as the constrained-expressions toolset. INCA Partial Evaluator AI-based Ada-to-SEDL State/event predicate definitions AI-variable bindings External events/states Promela True or Counter-example Ada Ada Source Ada System Completor Figure 4: Model Construction Process 6 EXPERIENCES In this section, we describe our experiences with applying the techniques described in this paper to model checking of a real partial software system. We begin with a description of this partial system. 6.1 Replicated Workers Computations The replicated workers framework (RWF) is a parameterizable parallel job scheduler, where the user configures the computation to be performed in each job, the degree of parallelism and several pre-defined variations of scheduler behavior. An instance of this framework is a collection of similar computational elements, called workers. Each worker repeatedly accesses data from a shared work pool, processes the data, and produces new data elements which are returned to the pool. User's define the number of workers, the type of work data, and computations to be performed by a worker on a data item. A version of this framework, written in Ada [15], implements workers, the pool and a lock as dynamically allocated instances of task types. Figure 5 illustrates the structure of the replicated workers framework and a sample of its interaction with a user application; procedure and entry calls are depicted with dashed and solid arrows, respectively, in the figure. Applications Create a collection of workers and a work pool and configure certain details of framework operation (e.g., whether the Execute routine operates as a Synchronous or Asynchronous invocation). A compu- Worker Worker Worker c := Create(.); Input(c, v); Create Input Execute User's Application doWork Figure 5: The Replicated Workers Framework tation is initialized through calls to the Input routine and started by calling Execute. Communicating only by way of the workpool, the collection of workers cooperate to perform the desired computation and terminate their execution when complete. Detailed description of the behavior of the RWF is provided in [15]. The execution state of the RWF consists of the local control flow states of a single pool, a single lock, and each of the workers. In addition, each of these tasks maintains local data. The original Ada code for the pool task, on the left of Figure 6, has a boolean variable, executeDone, three natural variables, numWait, numIdle, and workCount, two linked lists, WorkPool.List and newWork, two variables of the work type, and an array of task accesses workers, which is accessed through the discriminant value C. The lock task has a single boolean variable. Each worker task has a boolean variable, done, three linked lists, a task access variable, and an integer variable. In addition to the internal state of the RWF, we need to consider the state of the context, represented by stub and driver code. The only data component of that state is the work value passed to Input, which we refer to as driverInput. We will see, below, that these variables are abstracted in a variety of different ways in the finite-state models used for system validation. 6.2 Building RWF Models We use the approach described in Section 5 to produce finite-state systems that represent the behavior of the replicated workers framework. The frame-work is built of three active components: a task type (ActivePool) for the pool, a task type (ActiveWorker) for the worker, and a task which mediates access to a shared resource (ResultLock). The user is provided access to framework functionality through a collection of public procedures: a constructor Create, Input and task body ActivePool is Collection is a discriminant begin accept StartUp; workCount := 0; Outer: loop loop select * or accept Execute; C.done := FALSE; for i in 1 . C.max loop ? exit; * end select; loop select * or accept Put(newWork:in out WList) do Remove(newWork, workItem); for i in 1.Size(newWork) loop ?? Insert(work, workItem); workCount Remove(newWork, workItem); numIdle * end select; executeDone := TRUE; exit when Synchronous then accept Complete; C.done := TRUE; loop Outer; type ZERO POS is (zero, positive); task body GEN1ActivePool is choice begin accept StartUp; workCount := zero; Outer: loop loop select * or accept Execute; GEN1CollectionInfo done := FALSE; GEN2ActiveWorker.Execute; GEN3ActiveWorker.Execute; exit; * end select; loop select * or accept Put() do if choice then ?? workCount := positive; * end select; if numIdle=3 and workCount=zero then executeDone := TRUE; exit when numWait=3; accept Complete; GEN1CollectionInfo done := TRUE; loop Outer; Figure Original Ada and Abstracted, Specialized Ada for ActivePool Task Output routines, and a routine to Execute the com- putation. ActiveWorkers call user provided functions (doWork,doResults) to perform subcomputations on given work and result data. In validating the RWF implementation we assume that only one task creates and accesses each instance of the framework. This means that a single driver can be used to complete the system model; if the assumption is relaxed we would incorporate multiple drivers and a parallel component. We illustrate the analysis of a configuration of the RWF with three workers and Synchronous execution semantics. We will reason about local correctness properties of this system that are either internal to the RWF or related to the semantics of the RWF's application interface. For this reason, the external alphabet is empty. The stub generated by the algorithm in Figure 3 in this case consists of choices among calls to the RWF procedures. Defining the generic parameters and parameters to the Create call to be consistent with these assumptions enables program specialization to eliminate a number of program variables. In particular, the pool's work variables and array of task accesses, and each worker's three linked lists, task access variable, and integer variable are eliminated. Some of these variables were eliminated because their values are known to be constant, others are eliminated, by copy propagation, because they only transfer values between other variables. A significant number of variables, ranging over large domains, remain in the program so we apply the AIs, from Section 4, to the remaining variables to construct three different abstracted versions of the RWF system. Model 1. This model is the most aggressively ab- stracted. The variables numWait, numIdle, workCount, WorkPool.List, newWork and driverInput are all abstracted to the point AI. Variables executeDone, done, and the lock's boolean retain their concrete semantics. Our initial attempts to validate RWF properties did not use this model; we used model 2. We developed this model in order to see if any of the existing properties could be checked on a more compact model than 2. The results presented below confirm that this was possible. Parameters passed to doWork and doResult routines, which are modeled with stubs, also require abstraction. In this model the input work parameter is abstracted with the point AI and the boolean output parameter uses the choice AI. Model 2. Figure 6 gives the original Ada source and the abstracted, specialized Ada code for the ActivePool task of the RWF. Due to space limitations some details of the example are elided from the Figure, denoted by *, but the most interesting transformations remain. As with the first model, several local vari- ables, WorkPool.List, newWork and driverInput, are abstracted to the point AI. Where those variables can influence branch decisions, non-determinism is used. Since there is no non-deterministic choice construct in Ada, we introduce a new variable choice that indicates, by convention, to the model construction tools that a non-deterministic choice of the value of the variable is desired. Of the remaining variables, executeDone, numWait and numIdle retain their concrete semantics and workCount is abstracted with a zero-pos AI. We note that that numWait and numIdle both act as bounded counters up to the number of workers, thus they have a relatively small impact on the size of the model. ActiveWorker tasks for this model are the same as for model 1. Some details of the specialization process are illustrated in Figure 6. Knowledge of the number of workers is exploited to unroll the ? loop and specialize its body. As a consequence of this, the resulting Ada contains only static task references (e.g., GEN1ActiveWorker). In fact, partial evaluation applied to this example converts all dynamically allocated data and tasks to a static form and all indirect data and task references to a static form. The ?? loop is not unrolled, rather, because of the zero-pos AI used in its body the specializer determines that there are only two possible values for workCount after the loop, unchanged and positive, and produces the conditional. Model 3. Model 2 was insufficient for validation of ordering properties of work items in the RWF. We constructed a third RWF model that incorporated the 2-ordered data AI for WorkType data and the 2-ordered list abstraction for the WorkPool.List data. Even though this model uses a non-trivial domain for variables of WorkType, the resulting model did not explicitly require the modeling of the pool's local variables, since they only serve to transfer values between lists. It is not possible to generate a compact, finite-state stub and driver that will input any sequence of work data to the partial system under analysis. For the properties we wish to check, such generality is not required of the stub and driver. In models 2 and 3 we require no information about the input sequence and consequently the point AI suffices. In this model, we require a finer abstraction. The stub and driver for this model incorporated the 2-ordered environment abstraction binding the driverInput variable with the 2-ordered data AI. Thus, input sequences are modeled as sequences of values from fd 1 ;d 2 ; otg. 6.3 The Properties We model checked a collection of correctness requirements of the replicated workers framework. The requirements were derived from an English language description of the framework and encoded as LTL formulae using patterns [12]. We expressed all the formulae in terms of event and state predicates that are converted automatically by the Inca toolset into propositions for use in defining LTL formulae for SPIN. An event refers to the occurrence of a rendezvous, a procedure call, or some other designated program statement. An event predicate is true if any task containing the specified event is in a state immediately following a transition on that event. State predicates define the points at which selected program variables hold a given value (e.g., states in which workCount is zero). We note the defining boolean expressions for encoding state predicates can be quite involved in some cases. For example, the Inca predicate definition for states where workCount is zero: (defpredicate "workCountISzero" (in-task activepool-task (= workCount "zero"))) causes the generation of a disjunction of 123 individual state descriptions, i.e., one for each state of the ActivePool task in which workCount has the value zero. Our experience suggests that automated support for such definitions is a necessary component of any finite-state software verification toolset. A selection of the specifications we checked are given in Figure 7. All model checks were performed using SPIN, version 3.09, on a SUN ULTRA5 with a 270Mhz UltraSparc IIi and 128Meg of RAM. Figure 8 gives the data for each of the model checking runs; the transition system model used for the run is given. We report the user+system time for running SPIN to convert LTL to the SPIN input format, to compile the Promela into a model checker, and to execute that model checker. The model construction tools were run on an AlphaStation 200 4/233 with 128Meg of RAM. The longest time taken (1) []((call doResults:i && !?return doResults:i) -? ((!call doResults:j) U return doResults:i)) Mutually exclusive execution of doResults. (2) (!?call Execute) -? ((!call doWork) U call Execute) No work is scheduled before execution. (3) []((return Execute && (!?call Execute)) -? ((!call doWork) U call Execute)) scheduled after termination. [](call Execute -? ((!return Execute) U (done w1 jj done w2 jj done w3 jj workCountEQZero jj [](!return Execute)))) Computation terminates when workpool is empty or worker signals termination. !(done w1 jj done w2 jj done w3) -? !?call doWork) If a worker is ready to get work, the workpool is not empty and the computation is not done, then work is scheduled. schedules work in input order. After a work item is scheduled, it will not be scheduled again. After a work item is processed, it will not be scheduled. Figure 7: LTL Specifications to convert completed Ada to SEDL was for model 3; it took 66.3 seconds. Generating Promela from the SEDL can vary due to differences in the predicate definitions required for different properties. The longest time taken for this step was also for model 3; it took 16.4 seconds. 6.4 Discussion All specified properties were known to hold on the RWF implementation we analyzed. For specifications 1-3 no filters were required to obtain true results. The remaining specifications required some form of filter. Properties 6-8 required the filter in Section 4 to insure the safety of the model check results under the 2-ordered AI incorporated in the transition system. We discuss the filters for properties 4-5 below. Specifications 1 and 7 are short-hand for collections of specifications for all are different worker tasks ids. The model checks times were equal for the different versions of each specification; one such time is given in Figure 8. Modeling missing components using the permissive stubs and drivers described in Section 5 has the advantage of yielding safe models for the system configuration considered. Its drawback is that it may not precisely describe the required behavior of missing components. This is the reason that model checks for specifications 4 Property Time Result Model (4f) 0.7, 3:44.3, 8.8 true 2 (6f) 1.3, 36:58.7, 13:08.1 true 3 Figure 8: Performance Data and 5 failed. To boost precision in analyzing these prop- erties, we code assumptions about the required behavior of missing components (e.g. doWork) as filter-formulae that are then model checked. Analysis of the counter example provided by SPIN for specification 4 showed that doResults calls made from GEN1ActiveWorker can call the ActivePool.Finished entry. This is because the stub routine allows doResults to perform any computation. API documentation for the RWF warns users against calling RWF operations from doWork and doResults. If we assume that users heed this warning, we can define two filters that eliminate such calls. The resulting filter for- mula, (4f), is: ([](call-Execute -? ((!return-Execute) U workCountISZero - [](!return-Execute))))) where call stubRWF w and call stubRWF r are rather large conjunctions of the events that correspond to the calls of RWF operations from stubs inlined at doWork and doResult call-sites within workers. The generation of these propositions is relatively simple using Inca's predicate definition mechanism. The same filters were used for specification (5f). The use of filters in properties (6-8f) was required since the AIs incorporated in the model were only guaranteed to be safe under the assumption of a single Insert of each work datum into the abstracted WorkPool.List. Even though the unfiltered versions of those properties returned a true result, those results cannot be trusted. It may be the case that the AI caused certain possible system executions to be excluded from the analysis. If such an execution violated the specified property then a true result might be returned when there is a defect in the system. To insure that this is not the case, we checked the following filter formula (6f): Properties (7-8f) only referred to d1 so they only include the filter that restricts Inserts of d1. 6.5 Lessons Learned Our experience using filters for model checking with this example is consistent with previous work on filter-based verification [11, 14, 28]. In many cases, no filters are necessary and when necessary relatively few filters are sufficient to achieve the level of precision necessary for property verification. Model checking of properties for our sample system was fast enough to be usable in a practical development setting. We note that the the second component of model check time in Figure 8 is the sum of the time for SPIN to compile Promela input to a C program and the time to compile that C program. The bulk of this cost in all cases was compiling the C program. The reader should not interpret these times as an inherent component of the cost of using SPIN. The Inca tools that we use to generate Promela code encode local task data into the control flow of a Promela task rather than as Promela variables. This can cause a dramatic increase in the size of the C program generated and consequently the compile times are significant. Further study is required to determine whether a more direct mapping to Promela would yield significant reductions in these times. It is interesting to note that the addition of filters can in some cases reduce analysis cost (e.g., property (4f)) while in others it can dramatically increase analysis cost (e.g., property (6f)). Conceptually, analysis cost can be reduced because paths through the finite-state model that are inconsistent with the filters are not considered during model checking. Analysis cost can be increased, on the other hand, because the effective state space (i.e., the product of the model and the property) is significantly larger. Further study is needed to understand the situations in which reduction or increase in analysis cost can be expected when using filters. Our methodology for incorporating AIs into finite-state models yields aggressively abstracted transition sys- tems. Nevertheless, as one might expect, even the relatively small changes in the abstractions we incorporated into our three models dramatically change the space, and consequently the time, required for model check- ing. Checks for properties (1) and (2), using model 1, required on the order of 1000 states to be searched, whereas checks for properties (6-8), using model 3, required on the order of 100000 states to be searched. We believe that constructing compact transition systems, while retaining sufficient precision to enable successful model checks, requires that AIs be selected and incorporated independently for properties that refer to the same set of propositions. In our experiments, the cost of model construction is not the dominant factor in analysis time, so the benefits of independently abstracted models may yield an overall reduction in analysis time. It is important to note that these observations are based on checking of local properties of a cohesive partial system with a relatively narrow and well-defined inter- In principle, it may be necessary to include a very large number of filters which may dramatically increase model check cost. We believe that application of the approach described in this paper is most sensible at points in the development process where unit and integration- level testing is currently applied. In this context, we believe that the sub-systems under analysis will be similar to the RWF system (i.e., highly cohesive and loosely coupled to the environment). Further evaluation is necessary to determine this conclusively and to study the cost of filter-based model checking for partial systems that are strongly coupled to their environment. The work described in this paper touches on model checking of software systems, model checking of open or partial systems and abstractions in model checking. In Section 3, we have already discussed the bulk of the related work. There have been some recent efforts to apply model checking techniques to abstracted software systems (e.g., [13, 30]). In that work, ad-hoc abstraction was performed by hand transforming source code into models suitable for analysis. While automating the selection of abstractions is a very difficult problem, application of abstractions is relatively well-understood. Unlike ad-hoc methods, our work builds off the solid semantic foundations and rich history of existing abstractions that have been developed in the twenty year history of abstract interpretation [10]. Furthermore, we explore the use of partial evaluation techniques (e.g., [21]) as a means of automating application of those abstractions. Our use of filters to refine a model of the environment is similar to other work on compositional verifica- tion. These divide-and-conquer approaches, decompose a system into sub-systems, derive interfaces that summarize the behavior of each sub-system (e.g., [6]), then perform analyses using interfaces in place of the details of the sub-systems. This notion of capturing environment behavior with interfaces also appears in recent developments on theoretical issues related to model checking of partial systems (e.g., [22, 23]). There has been considerably less work on the practical issues involved with finite-state verification of partial systems. Aside from our work with FLAVERS, discussed in Section 3, there are two other recent related practical efforts. Avrunin, Dillon and Corbett [5] have developed a technique that allows partial systems to be described in a mixture of source code and specifications. In there work, specifications can be thought of as assumptions or filters on a naive completion of a partial system given in code. Unlike our work, their approach is targeted to automated analysis of timing properties of systems. Colby, Godefroid and Jagadeesan [8] describe an automatable approach to completing reactive partial sys- tem. Unlike our approach, there work is aimed at producing a completed system that is executable in the context of the VeriSoft toolset [17]. Their system completion acts as a controlling environment that causes the given partial system to systematically explore its behavior and compare it to specifications of correctness prop- erties. To produce a tractable completion, they perform a number of analyses to determine which portions of the partial system can be influenced by external behavior, for example, tests of externally defined variables are modeled with non-deterministic choice. This is equivalent to abstracting all external data with a point AI, which happens by default in our approach. Our use of filters allows restriction of external component behavior which is not possible in their approach. Both approaches are sensitive to elimination, or abstraction, of program actions that may cause run-time errors; in our case this is manifested by modeling self-entry calls as error actions. We have described an automatable approach to safely completing the definition of a partial software system. We have shown how a completed system can be selectively abstracted and transformed into a finite-state system that can be input to existing model checking tools. We have illustrated that this approach strikes a balance between size and precision in a way that enables model checking of system requirements of real software components. Finally, we have shown how to refine the representation of system behavior, in cases where the precision of the base representation is insufficient, to enable proof of additional system requirements. There are a number of questions we plan to investigate as follow on work. In the work described in this pa- per, we encode filter information into the model check on-the-fly, an alternate method is to encode it directly into the finite-state system. We are currently comparing these two approaches in order to characterize the circumstances in which one approach is preferable to the other. In this paper, we consider individual abstrac- tions, encoded as AI, but we know of systems where the desired abstraction is a composition of two AIs. We are investigating the extent to which construction of such compositions can be automated. Finally, we are continuing development of the tools that make up our approach and we plan to evaluate their utility by applying them to additional real software systems. Along with a completed toolset, we plan to produce a library of abstractions that users can selectively apply to program variables. This will allow non-expert users to begin to experiment with model checking of source code for real component-based software systems. 9 ACKNOWLEDGEMENTS Thanks to John Hatcliff and Nanda Muhammad for helping to specialize versions of the replicated workers framework by hand. Thanks to James Corbett and George Avrunin for access to the Inca toolset. Special thanks to James for responding to our request for a predicate definition mechanism with an implementa- tion. Thanks to the anonymous referee who gave very detailed and useful comments on the paper. --R Abstract Interpretation of Declarative Languages. Analyzing partially-implemented real-time systems Checking subsystem safety properties in compositional reachability analysis. Automatic verification of finite-state concurrent systems using temporal logic specifications Automatically closing open reactive programs. Evaluating deadlock detection methods for concurrent software. Abstract interpretation: A unified lattice model for static analysis of programs by construction or approximation of fixpoints. Modular flow analysis for concurrent soft- ware Property specification patterns for finite-state verification Model checking graphical user interfaces using abstractions. Data flow analysis for verifying properties of concurrent programs. An application-independent concurrency skeleton in Ada-95 Limiting state explosion with filter-based refinement Model checking for programming languages using VeriSoft. Staging static analysis using abstraction-based program specializa- tion Automatically specializing software for finite-state verification The model checker SPIN. Partial Evaluation and Automatic Program Generation. Module checking. Module checking revis- ited The Temporal Logic of Reactive and Concurrent Systems: Specification. Synthesis of communicating processes from temporal logic specifications. Symbolic Model Checking. Advanced Compiler Design and Imple- mentation In transition from global to modular temporal reasoning about programs. Model checking software systems: A case study. Specifying interesting properties of programs in propositional temporal logics. --TR Automatic verification of finite-state concurrent systems using temporal logic specifications Abstract interpretation of declarative languages In transition from global to modular temporal reasoning about programs Automated Analysis of Concurrent Systems with the Constrained Expression Toolset The temporal logic of reactive and concurrent systems Partial evaluation and automatic program generation Data flow analysis for verifying properties of concurrent programs A concurrency analysis tool suite for Ada programs Model checking software systems Checking subsystem safety properties in compositional reachability analysis Model checking large software specifications An application-independent concurrency skeleton in Ada 95 Analyzing partially-implemented real-time systems The Model Checker SPIN Model checking for programming languages using VeriSoft Applying static analysis to software architectures Model checking graphical user interfaces using abstractions Automatically closing open reactive programs Advanced compiler design and implementation Property specification patterns for finite-state verification Synthesis of Communicating Processes from Temporal Logic Specifications Expressing interesting properties of programs in propositional temporal logic Abstract interpretation Symbolic Model Checking State-Based Model Checking of Event-Driven System Requirements Evaluating Deadlock Detection Methods for Concurrent Software Module Checking Modular flow analysis for concurrent software --CTR Oksana Tkachuk , Matthew B. Dwyer, Adapting side effects analysis for modular program model checking, ACM SIGSOFT Software Engineering Notes, v.28 n.5, September Frank Huch, Verification of Erlang programs using abstract interpretation and model checking, ACM SIGPLAN Notices, v.34 n.9, p.261-272, Sept. 1999 Matthew B. Dwyer , John Hatcliff , Roby Joehanes , Shawn Laubach , Corina S. Psreanu , Hongjun Zheng , Willem Visser, Tool-supported program abstraction for finite-state verification, Proceedings of the 23rd International Conference on Software Engineering, p.177-187, May 12-19, 2001, Toronto, Ontario, Canada Patrice Godefroid , Lalita J. Jagadeesan , Radha Jagadeesan , Konstantin Lufer, Automated systematic testing for constraint-based interactive services, ACM SIGSOFT Software Engineering Notes, v.25 n.6, p.40-49, Nov. 2000 John Penix , Willem Visser , Eric Engstrom , Aaron Larson , Nicholas Weininger, Verification of time partitioning in the DEOS scheduler kernel, Proceedings of the 22nd international conference on Software engineering, p.488-497, June 04-11, 2000, Limerick, Ireland G. J. Holzmann , M. H. Smith, An Automated Verification Method for Distributed Systems Software Based on Model Extraction, IEEE Transactions on Software Engineering, v.28 n.4, p.364-377, April 2002 Victor A. Braberman , Miguel Felder, Verification of real-time designs: combining scheduling theory with automatic formal verification, ACM SIGSOFT Software Engineering Notes, v.24 n.6, p.494-510, Nov. 1999 Ji Zhang , Betty H. C. Cheng, Specifying adaptation semantics, ACM SIGSOFT Software Engineering Notes, v.30 n.4, July 2005 John Hatcliff , Matthew B. Dwyer , Hongjun Zheng, Slicing Software for Model Construction, Higher-Order and Symbolic Computation, v.13 n.4, p.315-353, Dec. 1, 2000 Ji Zhang , Betty H. C. Cheng, Model-based development of dynamically adaptive software, Proceeding of the 28th international conference on Software engineering, May 20-28, 2006, Shanghai, China John Penix , Willem Visser , Seungjoon Park , Corina Pasareanu , Eric Engstrom , Aaron Larson , Nicholas Weininger, Verifying Time Partitioning in the DEOS Scheduling Kernel, Formal Methods in System Design, v.26 n.2, p.103-135, March 2005

filter-based analysis;assume-guarantee reasoning;software verification and validation;model checking

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Automated test data generation using an iterative relaxation method.

An important problem that arises in path oriented testing is the generation of test data that causes a program to follow a given path. In this paper, we present a novel program execution based approach using an iterative relaxation method to address the above problem. In this method, test data generation is initiated with an arbitrarily chosen input from a given domain. This input is then iteratively refined to obtain an input on which all the branch predicates on the given path evaluate to the desired outcome. In each iteration the program statements relevant to the evaluation of each branch predicate on the path are executed, and a set of linear constraints is derived. The constraints are then solved to obtain the increments for the input. These increments are added to the current input to obtain the input for the next iteration. The relaxation technique used in deriving the constraints provides feedback on the amount by which each input variable should be adjusted for the branches on the path to evaluate to the desired outcome.When the branch conditions on a path are linear functions of input variables, our technique either finds a solution for such paths in one iteration or it guarantees that the path is infeasible. In contrast, existing execution based approaches may require an unacceptably large number of iterations for relatively long paths because they consider only one input variable and one branch predicate at a time and use backtracking. When the branch conditions on a path are nonlinear functions of input variables, though it may take more then one iteration to derive a desired input, the set of constraints to be solved in each iteration is linear and is solved using Gaussian elimination. This makes our technique practical and suitable for automation.

Introduction Software testing is an important stage of software develop- ment. It provides a method to establish confidence in the reliability of software. It is a time consuming process and accounts for 50% of the cost of software development [10]. Given a program and a testing criteria, the generation of test data that satisfies the selected testing criteria is a very difficult problem. If test data for a given testing criteria for a program can be generated automatically, it can relieve the software testing team of a tedious and difficult task, reducing the cost of the software testing significantly. Several approaches for automated test data generation have been proposed in the literature, including random [2], syntax based [5], program specification based [1, 9, 12, 13], symbolic evaluation [4, 6] and program execution based [7, 8, 10, 11, 14] test data generation. A particular type of testing criteria is path coverage, which requires generating test data that causes the program execution to follow a given path. Generating test data for a given program path is a difficult task posing many complex problems [4]. Symbolic evaluation [4, 6] and program execution based approaches [7, 10, 14] have been proposed for generating test data for a given path. In general, symbolic evaluation of statements along a path requires complex algebraic manipulations and has difficulty in handling arrays and pointer references. Program execution based approaches can handle arrays and pointer references efficiently because array indexes and pointer addresses are known at each step of program execution. But, one of the major challenges to these methods is the impact of infeasible paths. Since there is no concept of inconsistent constraints in these methods, a large number of iterations can be performed before the search for input is abandoned for an infeasible path. Existing program execution based methods [7, 10] use function minimization search algorithms to locate the values of input variables for which the selected path is traversed. They consider one branch predicate and one input variable at a time and use backtracking. Therefore, even when the branch conditions on the path are linear functions of input, they may require a large number of iterations for long paths. In this paper, we present a new program execution based approach to generate test data for a given path. It is a novel approach based on a relaxation technique for iteratively refining an arbitrarily chosen input. The relaxation technique is used in numerical analysis to improve upon an approximate solution of an equation representing the roots of a function [15]. In this technique, the function is evaluated at the approximate solution and the resulting value is used to provide feedback on the amount by which the values in the approximate solution should be adjusted so that it becomes an exact solution of the equation. If the function is lin- this technique derives an exact solution of the equation from an approximate solution in one iteration. For nonlinear functions it may take more than one iteration to derive an exact solution from an approximate solution. In our method, test data generation for a given path in a program is initiated with an arbitrarily chosen input from a given domain. If the path is not traversed when the program is executed on this input, then the input is iteratively refined using the relaxation technique to obtain a new input that results in the traversal of the path. To apply the relaxation technique to the test data generation problem, we view each branch condition on the given path as a function of input variables and derive two representations for this function. One representation is in the form of a subset of input and assignment statements along the given path that must be executed in order to evaluate the function. This representation is computed as a slicing operation on the data dependency graph of the program statements on the path, starting at the predicate under consideration. Therefore, we refer to it as a predicate slice. Note that a predicate slice always provides an exact representation of the function computed by a branch condition. Using this exact representation in the form of program statements, we derive a linear arithmetic representation of the function computed by the branch condition in terms of input variables. An arithmetic representation of the function in terms of input variables is necessary to enable the application of numerical analysis techniques since a program representation of the function is not suitable for this purpose. If the function computed by a branch condition is a linear function of the input, then its linear arithmetic representation is exact. When the function computed by a branch condition is a nonlinear function of the input, its linear arithmetic representation approximates the function in the neighborhood of the current input. These two representations are used to refine an arbitrarily chosen input to obtain the desired input as follows. Let us assume that by executing a predicate slice using the arbitrarily chosen input, we determine that a branch condition does not evaluate to the desired outcome. In this case, the evaluation of the branch condition also provides us with a value called the predicate residual which is the amount by which the function value must change in order to achieve the desired branch outcome. Now using the linear arithmetic representation and the predicate residual, we derive a linear constraint on the increments for the current input. One such constraint is derived for each branch condition on the path. These linear constraints are then solved simultaneously using Gaussian elimination to compute the increments for the current input. A new input is obtained by adding these increments to the current input. Since the constraints corresponding to all the branch conditions on the path are solved simultaneously, our method attempts to change the current input so that all the branch predicates on the path evaluate to their desired outcomes when their predicate slices are executed on the new input. If all the branch conditions on the path are linear functions of the input (i.e., the linear arithmetic representations of the predicate functions are exact), then our method either derives a desired input in one iteration or guarantees that the path is infeasible. This result has immense practical importance in accordance with the studies reported in [6]. A case study of 3600 test case constraints generated for a group of Fortran programs has shown that the constraints are almost always linear. For this large class of paths our method is able to detect infeasibility, even though the problem of detecting infeasible paths is unsolvable in general. If such a path is feasible, our method is extremely efficient as it finds a solution in exactly one iteration. If at least one branch condition on the path is a nonlinear function of the input, then the increments for the current input that are computed by solving the linear constraints on the increments may not immediately yield a desired input. This is because the set of linear constraints on the increments are derived from the linear arithmetic representations (which in this case are approximate) of the corresponding branch conditions. Therefore it may take more than one iteration to obtain a desired input. Even when the branch predicates on the path are nonlinear functions of the input, the set of equations to be solved to obtain a new input from the current input are linear and are solved by Gaussian elim- ination. Gauss elimination algorithm is widely implemented and is an established method for solving a system of linear equations. This makes our technique practical and suitable for automation. The important contributions of the novel method presented in this paper are: ffl It is an innovative use of the traditional relaxation technique for test data generation. ffl If all the conditionals on the path are linear functions of the input, it either generates the test data in one iteration or guarantees that the path is infeasible. There- fore, it is efficient in finding a solution as well as powerful in detecting infeasibility for a large class of paths. ffl It is a general technique and can generate test data even if conditionals on the given path are nonlinear functions of the input. In this case also, the number of iterations with inconsistent constraints can be used as an indication of a potential infeasible path. ffl The set of constraints to be solved in this method is always linear even though the path may involve conditionals that are non-linear functions of the input. A set of linear constraints can be automatically solved using Gaussian elimination whereas no direct method exits to solve a set of arbitrary nonlinear constraints. Gaussian elimination has been widely implemented and experimented algorithm. This makes the method practical and suitable for automation. ffl It is scalable to large programs. The number of program executions required in each iteration are independent of the path length and are bounded by number of input variables. The size of the system of linear equations to be solved using Gaussian elimination increases with the number of branch predicates on the path, but the increase in cost is significantly less than that of the existing techniques. The organization of this paper is as follows. An overview of the method is presented in the next section. The algorithm for test data generation is described in section 3. It is illustrated with examples involving linear and nonlinear paths, loops and arrays. Related work is discussed in section 4. The important features of the method are summarized and our future work is outlined in section 5. Overview We define a program module M as a directed graph e), where N is a set of nodes, E is a set of edges, s is a unique entry node and e is a unique exit node of M . A node n represents a single statement or a conditional expression, and a possible transfer of control from node n i to node n j is mapped to an edge (n G is a sequence of nodes such that (n A variable i k is an input variable of the module M if it either appears in an input statement of M or is an input parameter of M . The domain Dk of input variable i k is the set of all possible values it can hold. An input vector is the number of inputs, is called a Program Input. In this paper, we may refer to the program input by input and use these terms interchangeably. A conditional expression in a multi-way decision statement is called a Branch Predicate. Without loss of gen- erality, we assume that the branch predicates are simple relational expressions (inequalities and equalities) of the form E1 op E2 , where E1 and E2 are arithmetic expressions, and op is one of f!;-;=; 6=g. If a branch predicate contains boolean variables, we represent the "true" value of the boolean variable by a numeric value zero or greater and the "false" value by a negative numeric value. If a branch predicate on a path is a conjunction of two or more boolean variables such as in then such a predicate is considered as multiple branch predicates must simultaneously be satisfied for the traversal of the path. If a branch predicate on a path is a disjunction of two or more boolean variables such as in B), then at a time only one of the branch predicates A - 0 or B - 0 is considered along with other branch predicates on the path. If a solution is not found with one branch predicate then the other one is tried. Each branch predicate E1 op E2 can be transformed to the equivalent branch predicate of the form F op 0, where F is an arithmetic expression E1 \Gamma E2 . Along a given path, F represents a real valued function called a Predicate Function. F may be a direct or indirect function of the input variables. To illustrate this, let us consider the branch predicate for the conditional statement P2 in the example program in Figure 1. The predicate function F2 corresponding to the branch predicate BP2 is Along path g, the predicate function F2 indirectly represents the function of the input variables X;Y; Z. We now state the problem being addressed in this paper: Problem Statement: Given a program path P which is traversed for certain evaluations (true or false) of branch predicates BP1 a program input I that causes the branch predicates to evaluate such that P is traversed. We present a new method for generating a program input such that a given path in a program is traversed when the program is executed using this input. In this method, test data generation is initiated with an arbitrarily chosen input from a given domain. If the given path is not traversed on this input, a set of linear constraints on increments to the input are derived using a relaxation method. The increments obtained by solving these constraints are added to the input to obtain a new input. If the path is traversed on the new input then the method terminates. Otherwise, the steps of refining the input are carried out iteratively to obtain the desired input. We now briefly review the relaxation technique as used in numerical analysis for refining an approximation to the solution of a linear equation. The Relaxation Technique be an approximation to a solution of the linear equation In general, substituting (x0 ; y0) in the lhs of the above equation would result in a non zero value r0 called the residual, i.e., If increments \Deltax and \Deltay for x0 and y0 are computed such that they satisfy the linear constraint given by Therefore, is a solution of equation (1). In order to formulate the test data generation problem as a relaxation technique problem, we view the predicate function corresponding to each branch predicate on the path as a function of program input. To apply the above relaxation technique, a Linear Arithmetic Representation in terms of the relevant input variables is required for each predicate function. We first derive an exact program representation called a Predicate Slice for the function computed by each predicate function and then use it to derive a linear arithmetic representation. The two representations are used in an innovative way to refine the program input. The Predicate Slice The exact program representation of a predicate function, the Predicate Slice, is defined as follows: Definition: The Predicate Slice S(BP;P ) of a branch predicate BP on a path P is the set of statements that compute values upon which the value of BP may be directly or indirectly data dependent when execution follows the path P . In other words, S(BP;P ) is a slice over data dependencies of the branch predicate BP using a program consisting of 0: read(X,Y,Z) 1: 2: 3: else 5: 7: 8: write("Nonlinear: Quadratic") endif 9: endif 0: read(X,Y,Z) Statements in Predicate Slice S(BP1, P) 0: read(X,Y,Z) 1: 2: Statements in Predicate Slice S(BP2, P) 0: read(X,Y,Z) 1: Statements in Predicate Slice S(BP4, P) Figure 1: An Example Program and Predicate Slices on a path P=f0,1,P1,2,P2,4,5,6,P4,9g only input and assignment statements preceding BP on the path P . We illustrate the above definition using the example program in Figure 1. Consider the path P Let BP i denote the i th branch predicate along the path P . The predicate slices corresponding to the branch predicates BP 1, BP2 and BP4 along path P are: As illustrated by the above examples, predicate slices include only input and assignment statements. The value of a predicate function for an input can be computed by executing the corresponding predicate slice on the input. Note that a predicate slice is not a conventional static slice since it is computed over the statements along a path. It is also not a dynamic slice because it is computed statically using the input and assignment statements along a path and is not as precise as the dynamic slice. To illustrate the latter we consider the code segment given in Figure 2: input(I, J, Y); Figure 2: A code segment on a path using an array. When I 6= J, the evaluation of BP : (A[J] ? 0) is not data dependent on the assignment statement. Whereas, if I = J, the evaluation of BP is data dependent on the assignment state- ment. Therefore, the predicate slice for the branch predicate BP will consist of the input statement as well the assignment statement. In other words, the predicate slice is a path oriented static slice. The concept of predicate slice enables us to evaluate the outcome of each branch predicate on the path irrespective of the outcome of other branch predicates. The predicate slices for the branch predicates on the path can be executed using an arbitrary input even though the path may not be traversed on that input. This is possible because there are no conditionals in a predicate slice. After execution of a predicate slice on an input, the value of the corresponding predicate function can be computed and the branch outcome evaluated. There is a correspondence between the outcomes of the execution of the predicate slices on an input and the traversal of the the path on that input. If all the branch predicates on the path evaluate to their desired outcomes, by executing their respective predicate slices on an input and computing the respective predicate functions, the path will be traversed when the program is executed using this input. If any of the branch predicates on the path does not evaluate to its desired outcome when its predicate slice is executed on an input, the path will not be traversed when the program is executed using this input. Conceptually, a predicate slice enables us to view a predicate function on the path as an independent function of input variables. Therefore, our method can simultaneously force all branch predicates along the path to evaluate to their desired outcomes. In contrast, the existing program execution based methods [7, 10] for test data generation attempt to satisfy one branch predicate at a time and use backtracking to fix a predicate satisfied earlier while trying to satisfy a predicate that appears later on the path. They cannot consider all the branch predicates on the path simultaneously because the path may not be traversed on an intermediate input. The predicate slice is also useful in identifying the relevant subset of input variables, on which the value of the predicate function depends. This subset of input variables is required so that a linear arithmetic representation of the predicate function in terms of these input variables can be derived. The subset of the input variables on which the value computed by a predicate function depends can only be determined dynamically as illustrated by the example in Figure 2. Therefore, given an input and a branch predicate on the path, the corresponding predicate slice is executed using this input and a dynamic data dependence graph based upon the execution is constructed. The relevant input variables for the corresponding predicate function are determined by taking a dynamic slice over this dependence graph. Note that if only scalars are referenced in a predicate slice and the corresponding predicate function, then the subset of input variables on which the predicate function depends can be determined statically from the predicate slice. Execution of the predicate slice on the input data followed by a dynamic slice to determine relevant input variables is necessary to handle arrays. We define this subset of input variables as the Input Dependency Set. Definition: The Input Dependency Set ID(BP;I;P ) of a branch predicate BP on an input I along a path P is the subset of input variables on which BP is, directly or indirectly, data dependent. These input variables can be identified by executing the statements in the predicate slice S(BP;P ) on input I and taking a dynamic slice over the dynamic data dependence graph. For example, executing S(BP2;P ) on an input we note that the evaluation of BP2 depends on the input variables X,Y and Z. Therefore, Zg. Now we explain how we use the input dependency set to derive the linear arithmetic representation in terms of input variables for a predicate function for a given input. Deriving the Linear Arithmetic Representation of a Predicate Function Given a predicate function and its input dependency set ID for an input I, we write a general linear function of the input variables in ID. Then, we compute the values of the coefficients in the general linear function so that it represents the tangent plane to the predicate function at I. This gives us a Linear Arithmetic Representation for the predicate function at I. For example, the predicate function F2 has for the input linear function for the inputs in ID is Here, a; b and c are the slopes of f with respect to input variables X;Y and Z respectively and d is the constant term. If the slopes a, b and c above are computed by evaluating the corresponding derivatives of the predicate function at the input I0 and the constant term is computed such that the linear function f evaluates to the same value at I0 as that computed by executing the corresponding predicate slice on I0 and evaluating the predicate function, then represents the tangent plane to the predicate function at input I0 . This gives us the linear arithmetic representation for the predicate function at I0 . If the predicate function computes a linear function of the input, then the above tangent plane the exact representation for the predicate function. Whereas if a predicate function computes a nonlinear function of the input, then the above tangent plane f(X;Y; approximate the predicate function in the neighborhood of the input I0 . We illustrate this by deriving the linear arithmetic representation for the predicate function F2 at the input We approximate the derivatives of a predicate function by its divided differences. To compute a at I0 , we execute S(BP2;P ) at I0 and at where we have chosen increment in the input variable X. Then, we compute the divided difference: 2: This gives the value of a = 2: We compute the value of b by executing the predicate slice S(BP2;P ) at I0 and at and computing the divided difference of F2 at these two points with respect to y. This gives b equal to \Gamma2. Similarly, we get c equal to 1. We compute the value of d by solving for d from the equation a Substituting the values of a; b; c and F2(I0) in this equation and solving for d, we get d equal to \Gamma100. Therefore, we obtain the linear arithmetic representation for F2 at I0 as In this example, F2 computes a linear function of the input. Therefore, its linear arithmetic representation at I0 computed as above is the exact representation of the function of inputs computed by F 2. Also, only those input variables that influence the predicate function F2 appear in this representation In this paper, we have approximated the derivatives of a predicate function by its divided differences. Tools have been developed to compute derivative of a program with respect to an input variable [3]. With these tools, we can get exact derivative values rather than using divided differences. Therefore, our technique for deriving a linear arithmetic representation for a predicate function can be very accurately implemented for automated testing. Using the method explained above, we derive a linear arithmetic representation at the current input for each predicate function on the given path. In order to derive a set of linear constraints on the increments to the current input from these linear arithmetic representations, we execute the predicate slices of all the branch predicates on the current input and compute the values of corresponding predicate functions. We use these values of the predicate functions to provide feedback for computing the desired increments to the current input. The Predicate Residuals The values of the predicate functions at an input, defined as Predicate Residuals, essentially place constraints on the changes in the values of the input variables that, if satisfied, will provide us with a new input on which the desired path is followed. Definition: The Predicate Residual of a branch predicate for an input is the value of the corresponding predicate function computed by executing its predicate slice at the input. If a branch predicate has the relational operator then a non zero predicate residual gives the exact amount by which the value of the predicate function should change by modifying the input so that the branch evaluates to its desired outcome. Otherwise, a predicate residual gives the least (maximum) value by which the predicate function's value must be changed (can be allowed to change), by modifying the program input, such that the branch predicate evaluates (continues to evaluate) to the desired outcome. We explain this with examples given below. If a branch predicate evaluates to the desired outcome for a given input, then it should continue to evaluate to the desired outcome. In this case, the predicate residual gives the maximum value by which the predicate function's value can be allowed to change, by modifying the program input, such that the branch predicate continues to evaluate to the desired outcome. To illustrate this, let us consider the path P in the example program in Figure 1. Using an input the branch predicate BP2 evaluates to the desired branch for the path P to be traversed. The value of the predicate function F2 at I = (1; 2; 110) and hence the predicate residual at this input is 8. Therefore, the value of the predicate function can be allowed to decrease at most by 8 due to a change in the program input, so that the predicate function continues to evaluate to evaluate to a positive value. On the other hand, if a predicate does not evaluate to the desired outcome, the predicate residual gives the least value by which the predicate function's value must be changed, by modifying the program input, such that the branch predicate evaluates to the desired outcome. For example, using the input the branch predicate BP2 does not evaluate to the desired branch for the path P to be tra- versed. The value of the predicate function and hence the predicate residual at \Gamma99. Therefore, the input should be modified such that the value of the predicate function increases at least by 99 for the branch predicate BP2 to evaluate to its desired outcome. The predicate residuals essentially guide the search for a program input that will cause each branch predicate on the given path P to evaluate to its desired outcome. We compute a predicate residual at the current input for each branch predicate on the given path. Once we have a predicate residual and a linear arithmetic representation at the current input for each predicate function, we can apply the relaxation technique to refine the input. Refining the input The linear arithmetic representation and the predicate residual of a predicate function at an input essentially allow us to map the change in the value of the predicate function to changes in the program input. For each predicate function on the path P , we derive a linear constraint on the increments to the program input using the linear representation of the predicate function and the value of the corresponding predicate residual. This set of linear constraints is then solved simultaneously using Gaussian elimination to compute increments to the input. These increments are added to the input to obtain a new input. We illustrate the derivation of linear constraint corresponding to the predicate function F 2. The branch predicate BP2 evaluates to "false", when S(BP2; P ) is executed on the arbitrarily chosen input should evaluate to "true" for the path P to be traversed. The residual value \Gamma99 and the linear function are used to derive a linear constraint Note that the constant term d does not appear in this con- straint. Intuitively, this means that the increments to the input I0 should be such that the value of predicate function F2 changes more than 99 so as to force F2 to evaluate to a positive value and therefore force the corresponding branch predicate BP2 to evaluate to its desired outcome, i.e., "true" on the new input. For instance, is one of the solutions to the above constraint. We see that BP2 evaluates to "true" when S(BP2;P ) is executed on The linear constraint derived above from the predicate residual to compute the increments for the current input, is an important step of this method. It is through this constraint that the value of a predicate function at the current input provides feedback to the increments to be computed to derive a new input. Since this method computes a new program input from the previous input and the residuals, it is a relaxation method which iteratively refines the program input to obtain the desired solution. We would like to point out here that when the relational operator in each branch predicate on the path is "=", this method reduces to Newton's Method for iterative refinement of an approximation to a root of a system of nonlinear functions in several variables. To illustrate this, let us consider the linear constraint in equation (2). Let us assume that the relational operator in the corresponding branch predicate BP2 is "=" and for simplicity let F2 be a function of a single variable X. Then the linear constraint in equation 2 reduces to 99which is of the form In general, the branch predicates on a path will have equalities as well as inequalities. In such a case, our method is different from Newton's Method for computing a root of a system of nonlinear functions in several variables. But since the increments for input are computed by stepping along the tangent plane to the function at the current input, we expect our method to have convergence properties similar to Newton's Method. In our discussion so far, we have assumed that the conditionals are the only source of predicate functions. However, in practice some additional predicates should also be considered during test data generation. First, constraints on inputs may exist that may require the introduction of additional predicates (e.g., if an input variable I is required to have a positive value, then the predicate I ? 0 should be introduced). Second, we must introduce predicates that constrain input variables to have values that avoid execution errors (e.g., array bound checks and division by zero). By considering the above predicates together with the predicates from the conditionals on the path a desired input can be found. For simplicity, in the examples considered in this paper we only consider the predicates from the conditionals. 3 Description of the Algorithm In this section, we present an algorithm to generate test data for programs with numeric input, arrays, assignments, conditionals and loops. The technique can be extended to nonnumeric input such as characters and strings by providing mappings between numeric and nonnumeric values. The main steps of our algorithm are outlined in Figure 3. We now describe the steps of our algorithm in detail and at the same time illustrate each step of the algorithm by generating test data for a path along which the predicate functions are linear functions of the input. Examples with nonlinear predicate functions are given in the next section. The method begins with the given path P and an arbitrarily chosen input I0 in the input domain of the program. The program is executed on I0 . If P is traversed using I0 , then I0 is the desired program input and the algorithm ter- minates; otherwise the steps of iterative refinement using the relaxation technique are executed. We illustrate the algorithm using the example from section 2, where the path in the program of Figure 1, with initial program input considered. The path P is not traversed when the program is executed using I0 . Thus, the iterative relaxation method as discussed below is employed to refine the input. Step 1. Computation of Predicate Slices. For each node n i in P that represents a branch predicate, we compute its Predicate Slice S(n by a backward pass over the static data dependency graph of input and assignment statements along the path P before n i . The predicate slices for the branch predicates on the path P are: Step 2. Identifying the Input Dependency Sets. For every node n i in P that represents a branch predicate, we identify the input dependency set ID(n variables on which n i is data dependent by executing the predicate slice S(n on the current input Ik and taking a dynamic slice over the dynamic data dependence graph. The input dependency sets for the branch predicates on the path P computed by executing the respective predicate slices on P using the input are: Note that all input and assignment statements along the path P need be executed at most once to compute the input dependency sets for all the branch predicates along the path Step 3. Derivation of Linear Arithmetic Representations of the Predicate Functions. In this step, we construct a linear arithmetic representation for the predicate function corresponding to each branch predicate on P . For each branch predicate n i in P , we first formulate a general linear function of the input variables in the set ID(n For example, the linear formulations for the predicate functions corresponding to the branch predicates on path P are: The coefficients a i , b i and c i of the input variables in the above linear functions represent the slopes of the i th predicate function with respect to input variables X;Y and Z respectively. We approximate these slopes with respective divided differences. To compute the slope coefficient with respect to a variable, we execute the predicate slice S(n evaluate the predicate function F at the current input m) and at Ik the divided difference This gives the value of the coefficient of i j in the linear function for the predicate function F corresponding to node in P . Similarly, we compute the other slope coefficients in the linear function. In the example being considered, evaluating F1 by executing computing the divided difference with respect to X, we get larly, for F2 and F 4, we get 2: Computing the divided differences with respect to Y using (1; 2; computing the divided difference with respect to Z using (1; 2; and (1; 2; 4), we get To compute the constant term d i , we execute the corresponding predicate slice at Ik and evaluate the value of the predicate function. The values of input variables in Ik and the slope coefficients found above are substituted in the linear function, and it is equated to the value of the predicate function at Ik computed above. This gives a linear equation in one unknown and it can be solved for the value of the constant term. For the example being considered, we substitute the slope coefficients a i , b i and c i computed above and in the general linear formulations for the predicate functions F 1, F2 and F 4. Then, we equate the general linear formulations to their respective values at and obtain the following equations in d Solving for the constant terms d i , we get and Therefore, the linear arithmetic representations for the three predicate functions of P are given by: If a predicate function is a linear function of input variables then the slopes computed above are exact and this method results in the exact representation of the predicate function. Input: A path and an Initial Program Input I0 Output: A Program Input If on which P is traversed procedure TESTGEN(P;I0) If P traversed on I0 then step 1: for each Branch Predicate n i on P , do Compute while not Done do for each Branch Predicate n i on P , do step 2: Execute S(n input Ik to compute input dependency set ID(n step 3: Compute the linear representation L(ID(n for the predicate function for n i step 4: Compute step 5: Construct a linear constraint using R(n the computation of increment to Ik endfor step inequalities in the constraint set to equalities step 7: Solve this system of equations to compute increments for the current program input. Compute the new program input Ik+1 by adding the computed increments to Ik if the execution of the program on input Ik+1 follows path P then else k++ endif endwhile endprocedure Figure 3: Algorithm to generate test data using an iterative relaxation method. If a predicate function computes a nonlinear function, the linear function computed above represents the tangent plane to the predicate function at Ik . In the neighborhood of Ik , the inequality derived from the tangent plane closely approximates the branch predicate. Therefore, if the predicate function evaluates to a positive value at a program input in the neighborhood of Ik , then so does the linear function and vice versa. These linear representations and the predicate residuals computed in subsequent step are used to derive a set of linear constraints which are used to refine Ik and obtain Ik+1 . Step 4: Computation of Predicate Residuals. We execute the predicate slice corresponding to each branch predicate on P at the current program input Ik and evaluate the value of the predicate function. This value of the predicate function is the predicate residual value R(n the current program input Ik for a branch predicate n i on . The predicate residuals at I0 for the branch predicates on P are: Step 5: Construction of a System of Linear Constraints to be solved to obtain increments for the current input. In this step, we construct linear constraints for computing the increments \DeltaI k for the current input Ik , using the linear representations computed in step 3 and predicate residual values computed in step 4. We first convert the linear arithmetic representations of the predicate functions into a set of inequalities and equal- ities. If a branch predicate should evaluate to "true" for the given path to be traversed, the corresponding predicate function is converted into an inequality/equality with the same relational operator as in the branch predicate. On the other hand, if a branch predicate should evaluate to "false" for the given path, the corresponding predicate function is converted into an inequality with a reversal of the relational operator used in the branch predicate. If a branch predicate relational operator and should evaluate to "false" for the given path to be traversed, then we convert it into two inequalities, one with the relational operator ? and the other with the relational operator !. If the corresponding predicate function evaluates to a positive value at Ik , then we consider the inequality with ? operator else we consider the one with ! operator. This discussion also holds when a branch predicate has 6= relational operator and should evaluate to "true" for the given path to be traversed. This set of inequalities/equalities gives linear representations of the branch predicates on P as they should evaluate for the given path to be traversed. Converting the linear arithmetic representations for the predicate functions on the path P into inequalities, we get: Now using the linear arithmetic representations at Ik of the branch predicates as they should evaluate for the traversal of path P and using the predicate residuals computed at Ik , we apply the relaxation technique as described in the previous section to derive a set of linear constraints on the increments to the input. By applying the relaxation technique to the linear arithmetic representations computed above and the predicate residuals computed in the previous step, the set of linear constraints on increments to I0 are derived as given below: Note that the constant terms d i from the linear arithmetic representations do not appear in these constraints. Step Conversion of inequalities to equalities. In general, the set of linear constraints on increments derived in the previous step may be a mix of inequalities and equalities. For automating the method of computing the solution of this set of inequalities, we convert it into a system of equalities and solve it using Gaussian elimination. We convert inequalities into equalities by the addition of new variables. A simultaneous solution of this system of equations gives us the increments for Ik to obtain the next program input Ik+1 . Converting the inequalities to equalities in the constraint set, for the example being considered, by introducing three new variables u, v and w, we get: where we require that u; v; w ? 0. Step 7: Solution of the System of Linear Equations. We simultaneously solve the system of linear equations obtained in the previous step using Gaussian elimination. If the number of branch predicates on the path is equal to the number of unknowns (input variables and new variables) and it is a consistent nonsingular system of equations, then a straightforward application of Gaussian elimination gives the solution of this system of equations. If the number of branch predicates on the path is more than the number of unknowns, then the system of equations is overdetermined and there may or may not exist a solution depending on whether the system of equations is consistent or not. If the system of equations is consistent then again a solution can be found by applying Gaussian elimination to a subsystem with the number of constraints equal to the number of vari- ables, and verifying that the solution satisfies the remaining constraints. If the system of equations is not consistent, it is possible that the path is infeasible. In such a case, a consistent subsystem of the set of linear constraints is solved using Gaussian elimination and used as program input for the next iteration. A repeated occurrence of inconsistent systems of equations in subsequent iterations strengthens the possibility of the path being infeasible. A testing tool may choose to terminate the algorithm after a certain number of occurrences of inconsistent systems with the conclusion that the path is likely to be infeasible. If the number of branch predicates on the path is less than the number of unknowns, then the system of equations is underdetermined and there will be infinite number of solutions if the system is consistent. In this case, if there are n branch predicates, we select n unknowns and formulate the system of n equations expressed in these n unknowns. The other unknowns are the free variables. The n unknowns are selected such that the resulting system of equations is a set of n linearly independent equations. Then, this system of n equations in n unknowns is solved in terms of free variables, using Gaussian elimination. The values of free variables are chosen and the values of n dependent variables are com- puted. The solution obtained in this step gives the values by which the current program input Ik has to be incremented to obtain a next approximation Ik+1 for the program input. We execute the predicate slices and evaluate the predicate functions at the new program input Ik+1 . If all the branch predicates evaluate to their desired branches then Ik+1 is a solution to the test data generation problem. Otherwise, the algorithm goes back to step 2 with Ik+1 as the current program input for (K th iteration. As explained in the previous section, input dependency sets and the linear representations depend on the current input data. Therefore, they are computed again in the next iteration using Ik+1 . In the example considered, there are three linear constraints and six unknowns. Therefore, it is an underdetermined system of equations and can be considered as a system of three equations in three unknowns with the other three unknowns as free variables. If it is considered as a system of three equations in the three variables \DeltaX , \DeltaY and \DeltaZ and then Gaussian elimination is used to triangularize the coefficient matrix, we find that the third equation is dependent on first equation because the third row reduces to a row of zeros resulting in a singular matrix. Therefore, we consider it as a system of equations in \DeltaX , \DeltaZ and w:4 \DeltaX \DeltaZ The values of free variables can be chosen arbitrarily such that u, v and w ?0. Choosing the free variables u, v and \Deltay equal to 1, and solving for \Deltax, \Deltaz and w, we get, 2. Adding above increments to I0 , we get Executing the predicate slices on path P using input I1 and evaluating the corresponding predicate functions, we see that the three branch predicates evaluate to the desired branch leading to the traversal of P . Therefore, the algorithm terminates successfully in one iteration. In this method, a new approximation of the program input is obtained from the previous approximation and its residuals. Therefore, it falls in the class of relaxation meth- ods. The relaxation technique is used iteratively to obtain a new program input until all branch predicates evaluate to their desired outcomes by executing their corresponding predicate slices. If it is found that the method does not terminate in a given time, then it is possible that either the time allotted for test data generation was insufficient or there was an accumulation of round off errors during the Gauss elimination method due to the finite precision arithmetic used. Gaussian elimination is a well established method for solving a system of linear equations. Its implementations with several pivoting strategies are available to avoid the accumulation of round off errors due to finite precision arithmetic. Besides these two possibilities, the only other reason for the method not terminating in a given time is that the path is infeasible. It is clear from the construction of linear representations in step 3 that if the function of input computed by a predicate function is linear, then the corresponding linear representation constructed by our method is the exact representation of the function computed by the predicate function. We prove that in this case, the desired program input is obtained in only one iteration. Theorem 1 If the functions of input computed by all the predicate functions for a path P are linear, then either the desired program input for the traversal of the path P is obtained in one iteration or the path is guaranteed to be infeasible. Proof Let us assume that there are m input variables for the program containing the given path P and there are n branch predicates BP1;BP2; :::BPn on the path P such that n1 of them use use the relational ". The linear representations for the predicate functions corresponding to the predicates on P can be computed by the method described in Step 3 of the algorithm. Note that these representations will be exact because the functions of input computed by the predicate functions are linear. We can write the branch predicates on path P in terms of these representations as follows: eq. set 1 Note that the coefficients corresponding to input variables not in the input dependency set of a predicate function will be zero. be an approximation to the solution of above set of equations. Then we have: eq. set 2. where r i;j is the residual value obtained by executing the corresponding predicate slice using I0 and evaluating the corresponding predicate function. Let be a solution of the eq. set 1. Then, executing the given program at If would result in traversal of the path P . The goal is to compute this solution. Substituting If in eq. set 1, we get: eq. set 3 Now subtracting eq. set 2 from eq. set 3, we get: This is precisely the set of constraints on the increments to the input that must be satisfied so as to obtain the desired input. If the increment \Deltax i for x i0 is computed from the above set of constraints then gives the desired solution If . As illustrated above, this requires only one iteration. This indeed is the set of constraints used in Step 5 of our method for test data generation. Therefore, given any arbitrarily chosen input I0 in the program domain, our method derives the desired input in one iteration. While solving the constraints above, if it is found that the set is inconsistent then the given path P is infeasible. Therefore, our method either derives the desired solution in one iteration or guarantees that the given path is infeasible. 3.1 Paths with Nonlinear Predicate Slices. If the function of input computed by a predicate function is nonlinear, the predicate function is locally approximated by its tangent plane in the neighborhood of the current input Ik . The residual value computed at Ik provides feedback to the tangent plane at Ik for the computation of increments to Ik so that if the tangent plane was an exact representation of the predicate function, the predicate function will evaluate to the desired outcome in the next iteration. Because the slope correspondence between the tangent plane and the predicate function is local to the current iteration point Ik , it may take more than one iteration to compute a program input at which the execution of predicate slice results in evaluation of the branch predicate to the desired branch outcome. Let us now consider a path with a predicate function computing a second degree function of the input, for the example program in Figure 1, with initial program input The path P is not traversed on I0 . Therefore, input I0 is iteratively refined. The predicate slices and the input dependency sets of branch predicates BP1;BP2 and BP4 are the same as in the example on path with linear predicate slices. For BP 3, Also, the linear representations for the predicate functions F1; F2 and F4 are the same as for the example in the previous section. For F 3, we have The slope of F3 with respect to Z is computed by evaluating the divided difference at (1,2,3) and (1,2,4). The above linear function represents the tangent plane at I0 of the nonlinear function computed by the predicate function corresponding to branch predicate BP 3. Converting each of these functions into an inequality with the relational operator that the branch predicate should evaluate to, we get: Note that the relational operator for the representation for BP2 is different from that of the example in previous section because a different branch is taken. The predicate residuals at I0 for the predicate slices on P are: The set of linear constraints to be used for computing the increment for I0 using the results of above two steps are: 2: The inequalities in the above constraint set are converted to equalities by introducing new variables s ? resulting system of equations in \DeltaX , \DeltaY , \DeltaZ and v is solved using Gaussian elimination.61 \DeltaX \DeltaY \DeltaZ The free variables s, t and u are arbitrarily chosen equal to 1, and the system is solved for \DeltaX , \DeltaY , \DeltaZ , and v. The solution of the above system is: 2. These increments are added to I0 to obtain a new input I1 . Executing the predicate slices on P using the program input I1 , we find that all the four branch predicates evaluate to their desired branches resulting in the traversal of P . There- fore, the algorithm terminates successfully in one iteration. We summarize the results in the following table. Iteration This example illustrates that the tangent plane at the current input is a good enough linear approximation for the predicate function in the neighborhood of the current input. Now we consider a path with a predicate function computing the sine function of the input. Let us consider the following path P for the program in Figure 1, with initial program input The path P is not traversed on I0 because BP2 evaluates to an undesired branch on I0 . Therefore, the steps for iterative refinement of I0 are executed. We summarize the results of execution of our test data generation algorithm for this example in the table given below. For path P to be traversed, the branch predicates BP1 and BP4 should evaluate to "false" and the branch predicates BP2 and BP5 should evaluate to "true". As shown in the table, through iterations 1 to 4 of the algorithm BP 1, BP 2, and BP4 continue to evaluate to their desired outcomes and the values of inputs X, Y and Z are incremented such that F5 moves closer to zero in each iteration. Finally in iteration 4, F5 becomes positive for program input I4 and therefore BP5 becomes true. We would like to point out that if the linear arithmetic representation of a branch predicate is exact, then the branch predicate evaluates to its desired outcome in the first iteration and continues to do so in the subsequent it- erations. Whereas, if the linear arithmetic representation approximates the branch predicate in the neighborhood of current input (as in the case of BP5) by its tangent plane, then although in each iteration the refined input evaluates to the desired outcome with respect to the tangent plane, it may take several iterative refinements of the input for the corresponding branch predicate to evaluate to its desired outcome. In this example, BP5 evaluated to "true"(its desired out- come) at I0 , but it evaluated to "false" at I1 . This is because the predicate residual provides the feedback to the linear representation (i.e., the tangent plane to the sine function) of BP5 and the input is modified by stepping along this linear arithmetic representation. As a result, the linear representation evaluates to a positive value at I1 , but the change in the program input still falls short of making the predicate function F5 evaluate to a positive value at I1 . In the subsequent iterations, the input gets further refined and finally in the fourth iteration, the predicate function F5 evaluates to a positive value. As illustrated by this example, after the first iteration, all the branch predicates computing linear functions of the input continue to evaluate to their desired outcomes as the input is further refined to satisfy the branch predicates computing nonlinear functions of input. During regression test- ing, a branch predicate or a statement on the given path may be changed. To generate test data so that the modified path is traversed, an input on which other branch predicates already evaluate to their desired outcomes will be a good initial input to be refined by our method. Therefore, during regression testing, we can use the existing test data as the initial input and refine it to generate new test data. 3.2 Arrays and Loops When arrays are input in a procedure, one of the problems faced by a test data generation method is the size of the input arrays. Our test data generation method considers only those array elements that are referenced when the predicate slices for the branch predicates on the path are executed and the corresponding predicate functions are evaluated. The input dependency set for a given input identifies the input variables that are relevant for a predicate function. There- fore, the test data generation algorithm uses and refines only those array elements that are relevant. This makes the test data generation independent of the size of input arrays. We illustrate how our method handles arrays and loops by generating test data for a program from [10] given in Figure 4. We take the same path and initial input as in [10] so that we can compare the performance of the two program execution based test data generation methods. Therefore, where denotes the j th execution of the predicate P with initial program input: The path P is not traversed on I0 . Therefore, the steps for iterative refinement of I0 are executed. Let l, h, s denote low, high, step respectively, and then the predicate slices and input dependency sets of the branch predicates on P are: The linear representations for the predicate functions corresponding to the branch predicates on P are: y, h, The predicate residuals at I0 for the predicate functions of the branch predicates on P are: The set of linear constraints to be used for computing increment to I0 using the results of above two steps are: The above inequalities are converted to equalities by introducing seven new variables a, b, c, d, e, f and g. where a, d, f ? 0; and b, c, e and g - 0. Considering it a system of equations expressed in unknowns \Deltal, \Deltas, d, \Deltax, \Deltay, b and e, we get:66 664 \Deltal \Deltas d \Deltax \Deltay The unknowns and free variables are selected so as to obtain a nonsingular system of equations. The values of free variables can be chosen arbitrarily with the constraints that a, d, f ? 0; and b, c, e and g - 0. The values of free variables f , \Deltaz, \Deltah, and g are chosen as 1. The value of free variable a is 3 for integer arithmetic. Solving for the unknown variables using Gaussian elimination, we get: 3: The new input generated after the first iteration is: The input values of A[39]; A[51] and A[63] are copied into A[89]; A[91] and A[93] respectively and then the increments computed in this iteration are added to A[89]; A[91] and A[93], giving: This step is important because the increments computed in the current iteration have to be added to the input used in the current iteration but the resulting values have to be copied into the array elements to be used in the next itera- tion. Only elements A[89]; A[91]; and A[93] are relevant for the next iteration. By executing the predicate slices for P on the program input I1 and evaluating the corresponding predicate func- tions, we see that all the branch predicates evaluate to their desired branch outcomes resulting in the traversal of P . All the predicate functions corresponding to branch predicates on P compute linear functions of input. Therefore, as expected the algorithm terminates successfully in one itera- tion. We summarize the results of this example in the table in Figure 4. Korel in [10] obtains test data for the above path in 21 program executions, whereas our method finds a solution in only one iteration with only 8 program exe- cutions. One program execution is used in the beginning to test whether path P is traversed on I0 , six additional program executions are required for computation of all the slope computations for linear representations and one more program execution is required to test whether path P is traversed on I1 . If we choose the path the set of linear constraints obtained in step 3 will be in- consistent. Since, all the predicate functions for this path compute linear functions of input, from Theorem 1, we conclude that this path must be infeasible. It is easy to check that P is indeed an infeasible path. var A: array[1.100] of integer; min, max, i:integer; 1: input(low,high,step,A); 2: min := A[low]; 3: 4: i := low do 5: P3: if min ? A[i] then 7: 8: output(min,max); Iteration # low high step A[39] A[51] A[63] A[89] A[91] A[93] Iteration # BP11 BP21 BP31 BP12 BP22 BP32 BP13 Figure 4: An example using an array and a loop. 4 Related Work The most popular approach to automated test data generation has been through path oriented methods such as symbolic evaluation and actual program execution. One of the earliest systems to automatically generate test data using evaluation only for linear path constraints was described in [4]. It can detect infeasible paths with linear path constraints but is limited in its ability to handle array references that depend on program input. A more recent attempt at using symbolic evaluation for test data generation for fault based criteria is described in [6]. In this work, a test data generation system based on a collection of heuristics for solving a system of constraints is developed. The constraints derived are often imprecise, resulting in an approximate solution on which the path may not be traversed. Since the test data is not refined further so as to eventually obtain the desired input, the method fails when the path is not traversed on the approximate solution. A program execution based approach that requires a partial solution to test data generation problem to be computed by hand using values of integer input variables is described in [14]. There is no indication on how to automate the step requiring computation by hand. Program execution based approaches for automated test data generation have been described in [11, 8], but they have been developed for statement and branch testing criteria. An approach to automatically generate test data for a given path using the actual execution of the program is presented in [10]. Another program execution based approach that uses program instrumentation for test data generation for a given path has been reported in [7]. These approaches consider only one branch predicate and one input variable at a time and use backtracking. Therefore, they may require a large number of iterations even if all the branch conditionals along the path are linear functions of the input. If several conditionals on the selected path depend on common input variables, a lot of effort can be wasted in backtracking. They cannot consider all the branch predicates on the path simultaneously because the path may not be traversed on an intermediate input. The concept of predicate slice allows us to evaluate each branch predicate on the path independently on an intermediate input even though the path may not be traversed on this input. This makes our technique more efficient than other execution based methods. Our method is scalable to large programs since the number of program executions in each iteration is independent of the path length and at most equal to the number of input variables plus one. If there are m input variables, in each iteration, at most m executions of the input and assignment statements on the given path are required to compute the slope coefficients. The values of the predicate functions at input Ik are known from the previous iteration. One execution of the input and assignment statements on the path is required to test whether the path is traversed on Ik+1 . Our method uses Gaussian elimination to solve the system of linear equations, which is a well established and widely implemented technique to solve a system of linear equations. Therefore, our method is suitable for automa- tion. The size of the system of linear equations to be solved increases with an increase in the number of branch conditionals on a path, but the increase in cost in solving a larger system is significantly less than that of existing execution based methods. Conclusions In this paper, we have presented a new program execution based method, using well established mathematical tech- niques, to automatically generate test data for a given path. The method is an innovative application of the traditional relaxation technique used in numerical analysis to obtain an exact solution of an equation by iterative improvement of an approximate solution. The results obtained from this method for test data generation are very promising. It provides a practical solution to the automated test data generation problem. It is easy to implement as the tools required are already available. It is more efficient than existing program execution based approaches as it requires fewer program executions because all the branch predicates on the path are considered simultaneously, and there is no back- tracking. It can also detect infeasibility for a large class of paths in a single iteration. Because it is execution based, it can handle different programming language features. We are working on extending the technique for strings and pointers. --R "Test Plan Generation using Formal Grammars," "Automatic Generation of Random Self-checking Test "ADIC: An Extensible Automatic Differentiation Tool for ANSI-C," "A System to Generate Test Data and Symbolically Execute Programs," "A Rule Based Software Test Data Generator," "Constraint-based Automatic Test Data Generation," "ADTEST: A Test Data Generation Suite for Ada Software Systems," "Automatic Test Data Generation using Constraint Solving Techniques," "ATLAS - An Automated Software Testing System," "Automated Software Test Data Generation," A Dynamic Approach of Test Data Gener- ation "An Automatic Data Generation System for Data Base Simulation and Testing," "Automatic Generation of Testcase Datasets," "Automatic Generation of Floating-Point Test "Numerical Analysis," --TR Automatic generation of random self-checking test cases Automated Software Test Data Generation Constraint-Based Automatic Test Data Generation ADTEST Automatic test data generation using constraint solving techniques A Rule-Based Software Test Data Generator Test plan generation using formal grammars ATLAS-An Automated Software Testing System --CTR JinHui Shan , Ji Wang , ZhiChang Qi , JianPing Wu, Improved method to generate path-wise test data, Journal of Computer Science and Technology, v.18 n.2, p.235-240, March Nguyen Tran Sy , Yves Deville, Consistency techniques for interprocedural test data generation, ACM SIGSOFT Software Engineering Notes, v.28 n.5, September Jon Edvardsson , Mariam Kamkar, Analysis of the constraint solver in UNA based test data generation, ACM SIGSOFT Software Engineering Notes, v.26 n.5, Sept. 2001 Xun Yuan , Atif M. Memon, Using GUI Run-Time State as Feedback to Generate Test Cases, Proceedings of the 29th International Conference on Software Engineering, p.396-405, May 20-26, 2007 Stephen Thomas , Laurie Williams, Using Automated Fix Generation to Secure SQL Statements, Proceedings of the Third International Workshop on Software Engineering for Secure Systems, p.9, May 20-26, 2007 Meng-Luo Ji , Ji Wang , Shuhao Li , Zhi-Chang Qi, Automated WCET analysis based on program modes, Proceedings of the 2006 international workshop on Automation of software test, May 23-23, 2006, Shanghai, China Hari Hampapuram , Yue Yang , Manuvir Das, Symbolic path simulation in path-sensitive dataflow analysis, ACM SIGSOFT Software Engineering Notes, v.31 n.1, January 2006 Paolo Tonella, Evolutionary testing of classes, ACM SIGSOFT Software Engineering Notes, v.29 n.4, July 2004 Shan Lu , Pin Zhou , Wei Liu , Yuanyuan Zhou , Josep Torrellas, PathExpander: Architectural Support for Increasing the Path Coverage of Dynamic Bug Detection, Proceedings of the 39th Annual IEEE/ACM International Symposium on Microarchitecture, p.38-52, December 09-13, 2006 Torsten Robschink , Gregor Snelting, Efficient path conditions in dependence graphs, Proceedings of the 24th International Conference on Software Engineering, May 19-25, 2002, Orlando, Florida David Chays , Saikat Dan , Phyllis G. Frankl , Filippos I. Vokolos , Elaine J. Weber, A framework for testing database applications, ACM SIGSOFT Software Engineering Notes, v.25 n.5, p.147-157, Sept. 2000 Rupak Majumdar , Koushik Sen, Hybrid Concolic Testing, Proceedings of the 29th International Conference on Software Engineering, p.416-426, May 20-26, 2007 Wei Zhao , Lu Zhang , Yin Liu , Jiasu Sun , Fuqing Yang, SNIAFL: Towards a static noninteractive approach to feature location, ACM Transactions on Software Engineering and Methodology (TOSEM), v.15 n.2, p.195-226, April 2006 Gregor Snelting , Torsten Robschink , Jens Krinke, Efficient path conditions in dependence graphs for software safety analysis, ACM Transactions on Software Engineering and Methodology (TOSEM), v.15 n.4, p.410-457, October 2006 Cristian Cadar , Vijay Ganesh , Peter M. Pawlowski , David L. Dill , Dawson R. Engler, EXE: automatically generating inputs of death, Proceedings of the 13th ACM conference on Computer and communications security, October 30-November 03, 2006, Alexandria, Virginia, USA Marc Fisher, II , Gregg Rothermel , Darren Brown , Mingming Cao , Curtis Cook , Margaret Burnett, Integrating automated test generation into the WYSIWYT spreadsheet testing methodology, ACM Transactions on Software Engineering and Methodology (TOSEM), v.15 n.2, p.150-194, April 2006

predicate sliccs;input dependency set;predicate residuals;relaxation methods;dynamic test data generation;path testing

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Learning to Recognize Volcanoes on Venus.

Dramatic improvements in sensor and image acquisition technology have created a demand for automated tools that can aid in the analysis of large image databases. We describe the development of JARtool, a trainable software system that learns to recognize volcanoes in a large data set of Venusian imagery. A machine learning approach is used because it is much easier for geologists to identify examples of volcanoes in the imagery than it is to specify domain knowledge as a set of pixel-level constraints. This approach can also provide portability to other domains without the need for explicit reprogramming; the user simply supplies the system with a new set of training examples. We show how the development of such a system requires a completely different set of skills than are required for applying machine learning to toy world domains. This paper discusses important aspects of the application process not commonly encountered in the toy world, including obtaining labeled training data, the difficulties of working with pixel data, and the automatic extraction of higher-level features.

Introduction Detecting all occurrences of an object of interest in a set of images is a problem that arises in many domains, including industrial product inspection, military surveil- lance, medical diagnosis, astronomy, and planetary geology. Given the prevalence of this problem and the fact that continued improvements in image acquisition and storage technology will produce ever-larger collections of images, there is a clear need for algorithms and tools that can be used to locate automatically objects of interest within such data sets. The application discussed in this paper focuses on data from NASA/JPL's highly successful Magellan mission to Venus. The Magellan spacecraft was launched from Earth in May of 1989 with the objective of providing global synthetic aperture radar (SAR) mapping of the entire surface of Venus. In August of 1990 the spacecraft entered a polar elliptical orbit around Venus. Over the next four years Magellan returned more data than all previous planetary missions combined (Saunders et al., 1992), specifically, over 30,000 1024 \Theta 1024 pixel images covering 98% of the planet's surface. Although the scientific possibilities offered by this data set are numerous, the sheer volume of data is overwhelming the planetary geology research community. Automated or semi-automated tools are necessary if even a fraction of the data is to be analyzed (Burl et al., 1994a). 1.1. Scientific Importance Volcanism is the most widespread and important geologic phenomenon on Venus (Saun- ders et al., 1992), and thus is of particular interest to planetary geologists studying the planet. From previous low-resolution data, it has been estimated that there are on the order of one million small volcanoes (defined as less than 20 km in diameter) that will be visible in the Magellan imagery (Aubele and Slytua, 1990). Understanding the global distribution and clustering characteristics of the volcanoes is central to understanding the geologic evolution of the planet (Guest et al., 1992; Crumpler et al., 1997). Even a partial catalog including the size, location, and other relevant information about each volcano would enable more advanced scientific studies. Such a catalog could potentially provide the data necessary to answer basic questions about the geophysics of Venus, questions such as the relationship between volcanoes and local tectonic structure, the pattern of heat flow within the planet, and the mechanics of volcanic eruption. 1.2. Impact of an Automated System A catalog of large Venusian volcanoes (greater than 20 km in diameter) has been completed manually (Crumpler et al., 1997; Stofan et al., 1992). However, by optimistic estimates the time for a geologist to generate a comprehensive catalog of small volcanoes on Venus would be ten to twenty man-years. In fact, in our experiments we have found that humans typically become quite fatigued after labeling only 50-100 images over a few days. Thus, large-scale sustained cataloging by geologists is not realistic even if they had the time to devote this task. An automated system would provide many benefits, including the ability to maintain a uniform, objective standard throughout the catalog, thereby avoiding the subjectivity and drift common to human labelers (Cooke 1991; Poulton 1994). Even a partially automated system that functions as an "intelligent assistant" would have considerable impact. 1.3. Motivation for a Learning Approach There are two approaches one could follow for building an automated volcano cataloging system. The first would be to develop hand-coded, volcano-specific detectors based on a high-level description of the problem domain provided by human ex- perts. There are, however, a number of drawbacks to this method. Geologists are quite good at identifying examples of the objects of interest, but it is often difficult for them to identify precisely which characteristics of each volcano in the image led to its detection. High-level features, such as circularity or the presence of a summit pit, are difficult to translate into pixel-level constraints. Finally, visual recognition of localized objects is a problem that arises in many domains; using a hand-coded approach would require a significant amount of reprogramming for each new domain. The second approach is to use learning from examples. Since the geologists can identify examples of volcanoes with relative ease, their domain knowledge can be captured implicitly through the set of labeled examples. Using a learning algorithm, we can extract an appearance model from the examples and apply the model to find new (previously unseen) volcanoes. This approach can also provide portability since the user must merely supply a new set of training examples for each new problem domain-in principle no explicit reprogramming is required. 1.4. Related Work Most prior work on automated analysis of remotely sensed imagery has focused on two problems: (1) classification of homogeneous regions into vegetation or land-use types, e.g., (Richards, 1986) and (2) detection of man-made objects such as airports, roads, etc. The first technique is not applicable to the volcano detection problem, and the second is inappropriate because naturally-occurring objects (such as volcanoes) possess much greater variability in appearance than rigid man-made objects. Several prototype systems (Flickner et al. 1995; Pentland, Picard, and Sclaroff, 1996; Picard and Pentland, 1996) that permit querying by content have been developed in the computer vision community. In general, these systems rely on color histograms, regular textures, and boundary contours or they assume that objects are segmented and well-framed within the image. Since the small volcanoes in the Magellan imagery cannot be characterized by regular textures or boundaries, none of these approaches are directly applicable to the volcano cataloging problem. (For example, we found that the edge contrast and noise level in the SAR images did not permit reliable edge-detection.) In general, there has been relatively little work on the problem of finding natural objects in a cluttered background when the objects do not have well-defined edge or spectral characteristics. Hough transform methods were used for the detection of circular geologic features in SAR data (Cross, 1988; Skingley and Rye, 1987) but without great success. For the small volcano problem, Wiles and Forshaw (Wiles and Forshaw, 1993) proposed using matched filtering. However, as we will see in Section 3, this approach does not perform as well as the learning system described in this paper. Fayyad and colleagues (Fayyad et al., 1996) developed a system to catalog sky objects (stars and galaxies) using decision tree classification methods. For this domain, segmentation of objects from the background and conversion to a vector of feature measurements was straightforward. A good set of features had already been hand-defined by the astronomy community so most of the effort focused on optimizing classification performance. In contrast, for the Magellan images, separating the volcanoes from the background is quite difficult and there is not an established set of pixel-level features for volcanoes. 1.5. The JARtool System JARtool (JPL Adaptive Recognition Tool) is a trainable visual recognition system that we have developed in the context of the Magellan volcano problem. The basic system is illustrated in Figure 1. Through a graphical user interface (GUI), which is shown in Figure 2, a planetary scientist can examine images from the Magellan CD-ROMs and label examples in the images. The automated portion of the system consists of three components: focus of attention (FOA), feature extraction, and classification. Each of these components is trained for the specific problem of volcano detection through the examples provided by the scientist. The specific approach taken in JARtool is to use a matched filter derived from training examples to focus attention on regions of interest within the image. Principal components analysis (PCA) of the training examples provides a set of domain-specific features that map high-dimensional pixel data to a low-dimensional feature space. Supervised machine learning techniques are then applied to derive a mapping from the PCA features to classification labels (volcano or non-volcano). The PCA technique, which is also known as the discrete Karhunen-Loeve transform (Fuku- naga, 1990), has been used extensively in statistics, signal processing, and computer vision (Sirovich and Kirby, 1987; Turk and Pentland 1991; Moghaddam and Pent- land, 1995; Pentland et al., 1996) to provide dimensionality reduction. PCA seeks a lower-dimensional subspace that best represents the data. An alternate approach is linear discriminant analysis (LDA) (Duda and Hart, 1973; Swets and Weng, 1996), which seeks a subspace that maximizes the separability between classes. However, in the volcano context, the non-volcano class is so complex that LDA methods at the pixel-level do not work well. 1.6. Outline In Section 2 the JARtool system design process is described with an emphasis on the real-world issues that had to be addressed before standard "off-the-shelf" classification learning algorithms could be applied. In Section 3 we provide an empirical evaluation of our learning-based system and compare the performance with that of human experts. Section 4 indicates the current status of the project. In Section 5 we discuss the lessons learned from the project and how these application lessons could provide useful directions for future machine learning research. In Convolve image with matched filter and select regions with highest response Project each candidate region onto a bank of filters derived by principal components analysis filter Discriminate between volcanoes and non-volcanoes in projected feature space Figure 1. Overview of the JARtool system. Figure 2. In addition to the standard image browsing and labeling capabilities, the JARtool graphical user interface enables the user to learn models of an object and then look for novel instances of the object. The image displayed here is a 30km \Theta 30km region on Venus containing a number of small volcanoes. (See Figure 5 to find out where the volcanoes are located.) Section 6 we conclude with a summary of the main points of the article and indicate directions for future work. 2. System Design 2.1. Magellan Imagery Pettengill and colleagues (Pettengill et al., 1991) give a complete description of the Magellan synthetic aperture radar system and associated parameters. Here we bright dark strong backscatter weak backscatter near-range flank far-range flank RADAR Figure 3. Because of the topography, the near-range volcano flanks scatter more energy back to the radar and appear bright. In contrast, the far-range flanks scatter energy away and appear dark. focus only on how the imaging process affects the appearance of the volcanoes in the dataset. Figure 2 shows a 30km \Theta 30km area of Venus as imaged by Magellan. This area is located near lat . Illumination is from the lower left and the pixel spacing 1 is 60m. Observe that the larger volcanoes in the image have the classic radar signature one would expect based on the topography; that is, the side of the volcano closest to the radar (near-range) appears bright and the side away from the radar (far-range) appears dark. The reason is that the near-range side scatters more energy back to the sensor than the surrounding flat plains, while the far-range side scatters most of the energy off into space. The brightness of each pixel is proportional to the log of the returned energy, so volcanoes typically appear as a bright-dark pair within a circular planimetric outline. Near the center, there is often a visible summit pit that appears as a dark-bright pair since the radar energy backscatters strongly from the far-range rim. However, if the pit is too small relative to the image resolution, it may not appear at all or may appear just as a bright spot. A high-level illustration of the imaging process is given in Figure 3. These topography-induced features are the primary visual cues that geologists report using to locate volcanoes. However, there are a number of other, more subtle cues. The apparent brightness of an area in a radar image depends not only on the macroscopic topography but also on the surface roughness relative to the radar wavelength. Thus, if the flanks of a volcano have different roughness properties than the surrounding plains, the volcano may appear as a bright or dark circular area instead of as a bright-dark pair. Volcanoes may also appear as radial flow patterns, texture differences, or as disruptions of graben. (Graben are ridges or grooves in the planet surface, which appear as bright lines in the radar imagery-see Figure 2.) 2.2. Obtaining a Labeled Training Database Although the Magellan imagery of Venus is the highest resolution available, expert geologists are unable to determine with 100% certainty whether a particular image feature is indeed a volcano. This ambiguity is due to a variety of factors such as image resolution, signal-to-noise level, and difficulties associated with interpreting SAR data. For the same image, different geologists will produce different labelings, and even the same geologist may produce different labelings at different times. To help quantify this uncertainty, the geologists are asked to assign the training examples to subjective probability "categories." Based on extensive discussions with the geologists, five categories are used. almost certainly a volcano (p - 0.98); the image clearly shows a summit pit, a bright-dark pair, and a circular planimetric outline. Category probably a volcano (p - 0.80); the image shows only two of the three category 1 characteristics. Category 3 - possibly a volcano (p - 0.60); the image shows evidence of bright- dark flanks or a circular outline; summit pit may or may not be visible. Categroy 4 - a pit (p - 0.50); the image shows a visible pit but does not provide conclusive evidence for flanks or a circular outline. Category 5 - not a volcano (p - 0.0). The probability p attached to category i is interpreted as follows. Given that a geologist has assigned an image feature to category i, the probability that the feature is truly a volcano is approximately p i . Figure 4 shows some typical volcanoes from each category. The use of quantized probability bins to attach levels of certainty to subjective image labels is not new. The same approach is used routinely in the evaluation of radiographic image displays to generate subjective ROC (receiver operating characteristic) curves (Bunch, 1978; Chesters, 1992). A simple experiment was conducted to assess the variability in labeling between two planetary geologists, who will be referred to as A and B. Both of these geologists were members of the Volcanism Working Group of the Magellan science team and have extensive experience in studying Earth-based and planetary volcanism. They have published some of the standard reference works on Venus volcanism (Guest et al., 1992; Aubele and Slyuta, 1990; Crumpler et al., 1997). Each geologist separately labeled a set of four images known as HOM4. The labels were then compared using a simple spatial thresholding step to determine the correspondence between label events from the two geologists. simply refers to a labeler circling an image feature and assigning a subjective confidence label.) The resulting confusion matrix is given in Table 1. The (i; j)th element of the confusion matrix counts the number of times that labeler A assigned a visual feature to category i while labeler B assigned the same feature to category j. Two label events are considered to belong to the same visual feature, if they are within a few pixels of each other. The (i; 5) entries count the Volcanoes Category 1: Category 2: Category 3: Category 4: Figure 4. A selection of volcanoes from each of the confidence categories. Table 1. Confusion matrix of geologist A vs. geologist B on HOM4. geologist B Label 1 Label 2 Label 3 Label 4 Label 5 geologist A Label Label Label instances where labeler A provided label i, but labeler B did not provide any label (and vice versa for the (5; Entry (5,5) is not well-defined. If both labelers agreed completely (same location and label for all events), the confusion matrix would have only diagonal entries. In the case shown in Table 1, there is clearly substantial disagreement, as evidenced by the off-diagonal elements in the matrix. For example, label 3's are particularly noisy in both ``directions.'' Label 3's are actually noisier than label 4's because there is greater variability in the appearance of category 3's compared to category 4's (4's are simple pits, while 3's are less well-defined). About 50% of the objects assigned label 3 by either labeler are not detected at all by the other labeler. On the other hand, only about 30% of the objects assigned label 4 and 10% of the objects assigned label 1 by one labeler are missed by the other. The confusion matrix clearly illustrates that there is considerable ambiguity in small volcano identification, even among experts. Success for the task can only be measured in a relative manner. To evaluate performance, we treat one set of labels as ground truth and measure how well an algorithmic detector agrees with this set of reference labels. In this paper, reference labels 1-4 are all considered to be true volcanoes for the purpose of performance evaluation. An alternative "weighted" performance metric is discussed in (Burl et al., 1994b). We also measure how well human labelers agree with the reference labels. Ideally, an algorithm should provide the same consistency with the reference labels as the human experts. A consensus labeling generated by a group of geologists working together and discussing the merits of each image feature is often used as the reference, since in general this labeling will be a more faithful representation of the actual ground truth. A typical consensus labeling is shown in Figure 5. From the geologists' point of view, it is a useful achievement to detect most of the category 1's and 2's, as the category 3's and 4's would not be used as part of a conservative scientific analysis. 2.3. Focus of Attention The first component in the JARtool system is a focus of attention that is designed to take as input an image and produce as output a discrete list of candidate volcano locations. In principle, every pixel in the image could be treated as a candidate location; however, this is too expensive computationally. A better approach is to use the FOA to quickly exclude areas that are void of any volcanoes. Only local regions passing the FOA are given to subsequent (computationally more expensive) processes. Hence, the FOA should operate in an aggressive, low-miss- rate regime because any missed volcanoes at this stage will be lost for good. The rate of false positives (false alarms) from the FOA is not critical; these may still be rejected by later stages (classification). Given the constraints of the FOA (low miss rate and low computational cost), a reasonable approach is to use a matched filter, i.e., a linear filter that matches the signal one is trying to find. The matched filter is optimal for detecting a known signal in white (uncorrelated) Gaussian noise (Duda and Hart, 1973). Of course, the volcano problem does not quite satisfy these underlying assumptions. Specifi- cally, the set of observed volcanoes show structured variations due to size, type of volcano, etc. rather than "isotropic" variations implicit with a signal plus white noise model. Likewise, the clutter background cannot be properly modeled as white noise. Despite these caveats, we have found empirically that the following modified matched filtering procedure provides a reasonable focus of attention mechanism. pixel region around the i-th training volcano. There is some loss of information due to this windowing process (especially for larger volcanoes). However, in our experiments, the results have not been particularly sensitive to the value of k spoiled pixels) (Burl et al., 1996). This may indicate that most of the information is concentrated at the center of the volcano (for example, Figure 5. Consensus labeling of a Magellan SAR image of Venus. The labeling shows the size, location, and subjective uncertainty of each image feature. The dashed box corresponds to the subimage shown in Figure 2. the transition in shading and presence of a summit pit) or that the matched filter is not able to exploit the information from the periphery-both explanations probably have merit. Each k \Theta k region can be normalized with respect to the local average image brightness (DC level) and contrast as follows: ~ Figure 6. The matched filter displayed as a template (left) and as a surface plot (right). The matched filter captures many of the characteristics that planetary geologists report using when manually locating volcanoes. In particular, there is a bright central spot corresponding to the volcanic summit pit and left-to-right bright-dark shading. is the mean of the pixels in v i , oe i is their standard deviation, and 1 is a k \Theta k matrix of ones. This normalization is essential since there are large fluctuations in the DC and contrast between images and even between local areas of the same image. A modified matched filter f is constructed by averaging the normalized volcano examples from the training set. Figure 6 shows the resulting filter. Applying the matched filter to an image involves computing the normalized cross-correlation of f with each k \Theta k image patch. The cross-correlation can be computed efficiently using separable kernel methods to approximate the 2-D kernel f as a sum of 1-D outer products (Treitel and Shanks, 1971). High response values indicate that there is strong correlation between the filter and the image patch. Candidate volcano locations are determined by thresholding the response values and spatially aggregating any threshold crossings that are within a prescribed distance from each other (default distance = 4 pixels). Obviously one concern with such a simple FOA is that if the population of volcanoes contains significant subclasses then a single filter would not be expected to perform well. However, experiments with an alternative mechanism that uses clustering to find several matched filters has provided only limited improvement (Stough and Brodley, 1997). 2.4. Feature Extraction A region of interest (ROI) identified by the focus of attention algorithm can be viewed as a point in a k 2 -dimensional space by stringing the k \Theta k pixel values out into a long vector. Note, however, that there is a loss of spatial neighborhood information. Algorithms that treat the data in this form will not explicitly know that certain pixels were adjacent in the original image data. Also, given the small number of training examples relative to the dimensionality of the data, there is little hope of learning anything useful without additional constraints. Experimental results with a variety of feedforward neural network classification models verified this hypothesis (Baldi, 1994). The training data were often linearly separable in pixel space resulting in an underconstrained training procedure that allowed the model to essentially memorize the training data perfectly, but with poor generalization to unseen data. Thus, direct use of the pixels as input to a classification algorithm is not practical. To work around the small number of training examples, we make use of the fact that for visual data, there is additional prior information that helps constrain the problem. Specifically, there is reason to believe that the volcanoes "live" on a low-dimensional manifold embedded in -dimensional space. Although the manifold is certainly nonlinear, we make use of the principal components analysis (PCA) paradigm to approximate the manifold with a low-dimensional hyperplane. This approximation can be viewed as a mapping from the high-dimensional pixel space to a lower dimensional feature space in which the features consist of linear combinations of the pixel values. We have also experimented with clustering the training data in pixel space and applying PCA separately to each cluster. This extension yields an approximation to the manifold by a union of hyperplanes. (See Section 2.7 for additional discussion.) Before presenting a more detailed view of the PCA approach, we remark that PCA is not the only method available for linear feature extraction. The assumption behind PCA is that it is important to find features that represent the data. Other approaches, such as linear discriminant analysis (LDA), seek to find discriminative features that separate the classes. In the context of finding volcanoes, however, the "other" class is quite complex consisting of all patterns that are not volcanoes. Direct application of LDA in pixel space leads to poor results. Recently, a method was proposed that combines PCA and LDA to find "most discriminative features" (Swets and Weng, 1996). In this approach, PCA is used on the pooled set of examples (volcanoes and non-volcanoes) to project the pixel data to a lower dimensional feature space. LDA methods are then applied in the projected space. Effectively this amounts to using a "linear machine" classifier (Duda and Hart, 1973) in the space of principal components features. In Section 3 we demonstrate that by performing PCA on only the positive examples and allowing more complex classifiers in PCA space, the JARtool algorithm is able to outperform the method of Swets and Weng by a significant margin. PCA can be summarized as follows. The goal is to find a q-dimensional subspace such that the projected data is closest in L 2 norm (mean square error) to the original data. This subspace is spanned by the eigenvectors of the data covariance matrix having the highest corresponding eigenvalues. Often the full covariance matrix cannot be reliably estimated from the number of examples available, but the approximate highest eigenvalue basis vectors can be be computed using singular value decomposition (SVD). Each normalized training volcano is reshaped into a vector and placed as a column in an n \Theta m matrix X , where n is the number of pixels in an ROI m is the number of volcano examples. The SVD produces a factorization of X as follows: For notational convenience, we will assume m is less than n. Then in Equation 2, U is an n \Theta m matrix such that U T and diagonal with the elements on the diagonal (the singular values) in descending order, and V is m \Theta m with I m\Thetam . Notice that any column of X (equivalently, any ROI) can be written exactly as a linear combination of the columns of U . Furthermore, if the singular values decay quickly enough, then the columns of X can be closely approximated using linear combinations of only the first few columns of U . That is, the first few columns of U serve as an approximate basis for the entire set of examples in X . Thus, the best q-dimensional subspace on which to project is spanned by the first q columns of U . An ROI is projected into this q-dimensional feature space as follows: where x is the ROI reshaped as an n-dimensional vector of pixels, u i is the i-th column of U , and y is the q-dimensional vector of measured features. Figure 7-b shows the columns of U reshaped as ROIs. The templates are ordered according to singular value so that the upper left template corresponds to the maximum singular value. Notice that the first ten templates (top row) exhibit structure while the remainder appear very random. This suggests using a subspace of dimension - 10. The singular value decay shown in Figure 7-c also indicates that 6 to 10 features will be adequate to encode most of the information in the examples. Indeed, parameter sensitivity experiments, which are reported in (Burl et al., 1996) show that values of q in the range 4-15 yield similar overall performance. 2.5. Classification The FOA and feature extraction steps transform the original Magellan images into a discrete set of feature vectors that can be classified with "off-the-shelf" learning al- gorithms. The remaining step is to classify ROIs into volcano or non-volcano. FOA and feature learning are based exclusively on positive examples (volcanoes). The classifier could also be trained in this manner. However, there are arguments (Fuku- naga, 1990) showing that single-class classifiers are subject to considerable error even in relatively low dimensions because the location of the "other" distribution (a) Training Volcanoes (b) Principal Components (c) Singular Values 9010Feature Number Figure 7. (a) The collection of volcanoes used for feature synthesis. (b) The principal components derived from the examples. (c) The singular values indicate the importance of each of the features for representing the examples. is unknown. Experiments based on non-parametric density estimation of the volcano class verified this hypothesis: the method gave poorer performance than the two-class methods described below. The negative examples were not used in the FOA and feature learning steps due to the complexity of the non-volcano class. Nonetheless these steps provide substantial conditioning of the data. For example, the FOA centers objects within a k \Theta k window. The feature extraction step uses prior knowledge about visual data (i.e., the fact that certain object classes can be modeled by linear combinations of basis functions) to map the data to a lower-dimensional space in which there is an improved opportunity for learning a model that generalizes to unseen data. Hence, in the PCA space it is reasonable to use supervised two-class learning tech- niques. We have experimented with a variety of algorithms including quadratic (or Gaussian) classifiers, decision trees, linear discriminant analysis, nearest neighbors using Euclidean and spatially weighted distance measures (Turmon, 1996), tangent distance (Simard, le Cun, and Denker, 1993), kernel density estimation, Gaussian mixture models, and feedforward neural networks (Asker and Maclin, 1997b; Cherkauer, 1996). All of these methods (with the exception of linear discriminant analysis) yielded similar performance on an initial test set of images. We interpret this to mean that the critical system design choices were already made, specifically in the feature learning stage; the choice of classifier is of secondary importance. In the experiments reported in Section 3, the quadratic classifier is used as the default since it is optimal for Gaussian data and provides posterior probability estimates, which can be thresholded to vary the trade-off between detection and false alarm rate. Letting designate the volcano class and ! 2 the non-volcano class, we have the following from Bayes' rule: where y is the observed feature vector. For the quadratic classifier, the posterior probabilities are estimated by assuming the class-conditional densities are Gaussian where the statistics of each class (- i and \Sigma i ) are estimated from labeled training data. 2.6. Summary of the Training Procedure In summary, training consists of a three-step process based on the geologist-labeled images: 1. Construct the FOA detection filter from the volcanoes labeled in the training images. Apply the FOA to the training images and then use the "ground truth" labels to mark each candidate ROI as a volcano or non-volcano. 2. Determine principal component directions from the ROIs that were detected in step 1 and marked as volcanoes. 3. Estimate the parameters of a classifier from the labeled feature vectors obtained by projecting all of the training data detected in step one onto the PCA templates of step two. ROIs marked as true volcanoes in step one serve as the positive examples, while ROIs marked as non-volcanoes serve as the negative examples. Figure 8. Example volcanoes from four different clusters and their respective cluster centers. Each row represents a sample of volcanoes that have been clustered together using K-means. Comment: This training procedure contains some non-idealities. For example, the positive examples supplied to the classifier are the same examples used to derive the features (principal component directions). It would clearly be better if the classifier were to receive a disjoint set of positive training examples, but given the limited number of examples, we compromised on the procedure described above. 2.7. Extension to the Basic Algorithm One objection to the baseline approach presented thus far is that there are various subtypes of volcanoes, each with unique visual characteristics. One would not expect the (approximate) hyperplane assumption implicit in the PCA approach to hold across different volcano subtypes. This limitation could affect the algorithm's ability to generalize across different regions of the planet, and in fact in the experiments reported later (Section 3), we have observed that the baseline system performs significantly worse on heterogeneous sets of images selected from various areas of the planet. One solution we investigated involves using a combination of classifiers in which each classifier is trained to detect a different volcano subclass. The outputs from all the classifiers are then combined to produce a final classification. Subclasses of volcanoes are found automatically by clustering the raw pixel representation of the volcanoes in the training set using k-means (Duda and Hart, 1973). In Figure 8 we show the results of clustering the volcanoes into four classes. Each row corresponds to a cluster; the first column shows the cluster center, while the other columns show selected instances. For each cluster, principal components analysis is performed separately yielding a set of features (basis functions) specific to a subclass of volcanoes. A classifier is then trained for each subclass, and in the final step the predictions of all the classifiers are combined into one. Details of the method for combining classifiers are given in (Asker and Maclin, 1997a). Experimental results comparing the combined classifier approach with the baseline are given in Section 3. 3. Performance Evaluation Initial experiments were conducted using a small set of images called HOM4 (denoting a set of four images which were relatively homogeneous in appearance). The results from these experiments were used to provide feedback in the algorithm development process and also served to fix the values of miscellaneous parameters such as the ROI window size, FOA threshold, number of principal components, and so forth. Because of this feedback, however, performance on HOM4 cannot be considered as a fair test of the system (since in effect one is training on the test data). In addition, HOM4 did not include enough images to provide a thorough characterization of the system performance and generalization ability. After these initial experiments, the algorithm and all the miscellaneous parameters were frozen at specific values (listed in the Appendix). Based on empirical sensitivity studies (Burl et al., 1996), we believe the system is relatively insensitive to the exact values of these parameters. Note that "freezing" does not apply to parameters normally derived during learning such as the matched filter, principal components, or statistics used by the classifier. These are recalculated for each experiment from the stated set of training examples. Extensive tests were conducted on three large image sets (HOM38, HOM56, HET36). The naming convention for the image sets is to use HOM if the set is considered homogeneous (images from the same local region) and HET if the set is heterogeneous (images selected from various locations). The numerical suffix indicates the number of images in the data set. Note that the smallest of these datasets covers an area of 450km \Theta 450km. A summary of the experiments and image sets is given in Table 2. The number of volcanoes listed corresponds to the number of label events in the "ground-truth" reference list, i.e., each label event is counted as a volcano regardless of the assigned confidence. The main conclusion from these tests was that the baseline system performed well on homogoeneous sets in which all images were taken from the same region of the planet, but performed poorly on heterogeneous sets in which images were selected randomly from various locations on the planet. To better understand this difference in performance, we conducted a follow-up experiment using a small set of heterogeneous images HET5. Our initial hypothesis was that the discrepancy occurred because the volcanoes from different regions looked different. However, what we found was that "knowing the local volcanoes" was not nearly as important as knowing the local non-volcanoes. The argument used to arrive at this conclusion is somewhat subtle, but is explained in detail in Section 3.4. Table 2. Experiments and image sets used to evaluate system performance. Experiment Image Set #Volcanoes Description Initial HOM4 160 4 images from lat Testing Extended HOM38 480 38 images from lat Testing HOM56 230 56 images from lat HET36 670 36 images from various locations Follow-up HET5 131 5 images from various locations 3.1. ROC and FROC As explained in Section 2.2, we evaluate performance by measuring how well a detector (algorithmic or human) agrees with a set of reference labels. A "detection" occurs if the algorithm/human indicates the presence of an object at a location where a volcano exists according to the reference list. Similarly, a "false alarm" occurs if the algorithm/human indicates the presence of an object at a location where a volcano does not exist according to the reference list. Consider a system which produces a scalar quantity indicating detection confidence (e.g., the estimated posterior class probability). By comparing this scalar to a fixed threshold, one can estimate the number of detections and false alarms for that particular threshold. By varying the threshold one can estimate a sequence of detection/false-alarm points. The resulting curve is known as the receiver operating characteristic (ROC) curve (Green and Swets, 1966; MacMillan and Creelman, 1991; Spackman, 1989; Provost and Fawcett, 1997). The usual ROC curve plots the probability of detection versus the probability of false alarm. The probability of detection can be estimated by dividing the number of detections by the number of objects in the reference list. Estimating the probability of false alarm, however, is problematic since the number of possible false alarms in an image is not well-defined. A practical alternative is to use a "free-response" ROC (Chakraborty and Winter, 1990), which shows the probability of detection versus the number of false alarms (often normalized per image or per unit area). The FROC methodology is used in all of experiments reported in this in particular, the x-axis corresponds to the number of false alarms per square kilometer. The FROC shares many of the properties of the standard ROC 2 . For example, the best possible performance is in the upper left corner of the plot so an FROC curve that is everywhere above and to the left of another has better performance. The FROC curve is implicitly parameterized by the decision threshold, but in practice the geologist would fix this threshold thereby choosing a particular operating point on the curve. 3.2. Initial Experiments Experiments on HOM4 were performed using a generalized form of cross-validation in which three images were used for training and the remaining image was reserved for testing; the process was repeated four times so that each image served once as the test image. This type of training-testing procedure is common in image analysis problems (Kubat, Holte, and Matwin, 1998). The system output was scored relative to the consensus labeling with all subjective confidence categories treated as true volcanoes. The FROC performance curve is shown in Figure 9a. The horizontal dashed line across the top of the figure (labeled FOA=0.35) shows the best possible system performance using an FOA threshold of 0.35. (The line is not at 100% because the FOA misses some of the true volcanoes.) The performance points of two individual geologists are also shown in the figure. Geologist A is shown with the '*' symbol, while geologist B is shown with the '+'. Note that for these images the system performance (at an appropriately chosen operating point) is quite close to that of the individual geologists. The effect of using different operating points is shown in table form in Figure 10a. 3.3. Extended Performance Evaluation 3.3.1. Homogeneous Images Given the encouraging results on HOM4, we proceeded to test the system on larger images sets. The HOM4 images were part of a 7 \Theta 8 block of images comprising a full-resolution Magellan "data product." Within this block 14 images were blank due to a gap in the Magellan data acquisition process. The remaining 38 (56 minus 4 minus 14) images were designated as image set HOM38. Training and testing were performed using generalized cross-validation in which the set of images was partitioned into six groups or "folds." Two of the images did not contain any positive examples, so these were used only for training. The other 36 images were partitioned randomly into six groups of six with the constraint that each group should have approximately the same number of positive examples. Five folds were used for training and the remaining fold was used for testing; the process was repeated so that each fold served once as the test set. This leave-out-fold method was used rather than leave-out-image to reduce the run time. The FROC performance is shown in Figure 9b (solid line). Since we did not have consensus labeling available for the entire image set, the labels of geologist A were used as the reference. The '+' symbol shows the performance of geologist B (relative to A), while the 'o' symbol shows the performance of one of the non- geologist authors (Burl). The performance of the algorithm is similar to the HOM4 case except at higher false alarm rates where the HOM38 performance is lower by approximately 12%. The discrepancy is probably due to differences in the FOA performance. Note that for HOM4 the FOA asymptote is around 94%, while it is only at 83% for HOM38. For comparison FROC curves are plotted for two other methods. The dashed curve labeled "FOA" shows the performance that could be achieved by using only a matched filter but with a lower threshold. The combination of matched filter and classification yields better performance than the matched filter alone. (Matched filtering was proposed as a possible solution to the volcano-detection problem in (Wiles and Forshaw, 1993)). Also shown is the FROC for the discriminative Karhunen-Loeve (DKL) approach (Swets and Weng, 1996), which combines principal components and linear discriminant analysis. Observe that the JARtool approach provides significantly better performance (an increase in detection rate by 10 percentage points or more for a given false alarm rate). For the HOM38 experiments, the training images and test images were geographically close together. To test the system's generalization ability, another experiment was performed in which training was carried out on HOM4+HOM38 and testing was carried out on a geographically distinct set of homogeneous images HOM56. The images were from the same latitude as the training images and visually appeared to have similar terrain and volcano types. For this data set, reference labels were provided by one of the non-geologist authors (Burl). The baseline performance is shown as a solid curve in Figure 9c. The clustering extension to the baseline algorithm was also applied to the data. The corresponding FROC is shown with the dashed curve. The clustering approach appears to provide a slight improvement over the baseline algorithm, consistent with other results reported in (Asker and Maclin, 1997a). However, the baseline algorithm is still used in the fielded system because of its simplicity and shorter training time. (These factors are believed to outweigh the marginal loss in performance.) Images Finally, the system was evaluated under the most difficult conditions. A set of 36 images (HET36) was selected from random locations on the planet. These images contained significantly greater variety in appear- ance, noisiness, and scale than the previous image sets. Training was done on HOM4+HOM38. The system FROC performance (relative to consensus) is shown in Figure 9d, and the performance at selected operating points is shown in Figure 10d. Here the classifier performs much worse than on the more homogeneous data sets. For example at 0.001 false alarms/km 2 the detection performance is in the 35-40% range whereas for all homogeneous image sets, the detection rates were consistently in the 50-65% range. For the few images where we also have individual labels, the geologists' detection performance is roughly the same as it was on the homogeneous images. From these results it appears that human labelers are much more robust with respect to image inhomogeneity. 3.4. Follow-Up Analysis To better understand the decreased performance on heterogeneous images, we conducted follow-up experiments on a smaller set of images (HET5). Performance on this set was also poor, and our initial hypothesis was that the degradation occurred because the volcanoes were somehow different from image to image. To investigate this possibility, we performed experiments with two different training paradigms: (1) cross-validation in which one image was left out of the training set and (2) cross-validation in which one example was left out of the training set. The first method will be referred to as LOI for "leave-out-image"; the second method will be referred to as LOX for "leave-out-example." (a) HOM4 (b) HOM38 Alarms km Detection Rate Alarms km Detection Rate (c) HOM56 (d) HET36 Alarms km Detection Rate Alarms km Detection Rate Figure 9. FROC curves showing the performance of the baseline algorithm on four image sets. Each figure shows the trade-off between detection rate and the number of false alarms per area. Note that the algorithm performs considerably better on the homogeneous image sets (a,b,c) than on the heterogeneous set (d).The discrete symbols (*,+,O) in (a) and (b) show the performance of human labelers. Two nearest-neighbor classification algorithms were evaluated in addition to the baseline Gaussian classifier. The nearest-neighbor algorithms were applied directly to the pixel-space regions of interest (ROIs) identified by the FOA algo- rithm. To allow for some jitter in the alignment between ROIs, we used the peak cross-correlation value over a small spatial window as the similarity measure. One nearest-neighbor algorithm was the standard two-class type in which an unknown Operating point: OP1 OP2 OP3 OP4 OP5 OP6 Threshold: 0.75 0.80 0.85 0.90 0.95 0.99 Detected Category 1 (%) 88.9 88.9 86.1 80.6 72.2 63.9 Detected Category 2 (%) 89.7 89.7 86.2 79.3 72.4 65.5 Detected Volcanoes (%) 82.2 81.0 79.8 74.9 69.3 62.6 False Alarms per image 19.5 18.5 15.0 13.0 9.5 6.0 False Alarms per (b) HOM38 Operating point: OP1 OP2 OP3 OP4 OP5 OP6 Threshold: 0.75 0.80 0.85 0.90 0.95 0.99 Detected Category 1 (%) 92.0 92.0 88.0 84.0 84.0 76.0 Detected Category 2 (%) 80.8 78.2 75.6 71.8 65.4 50.0 Detected Volcanoes (%) 68.0 65.0 63.5 60.3 54.6 48.4 False Alarms per image 10.6 8.8 7.3 5.8 3.9 2.1 False Alarms per (c) HOM56 Operating point: OP1 OP2 OP3 OP4 OP5 OP6 Threshold: 0.75 0.80 0.85 0.90 0.95 0.99 Detected Category 1 (%) 100.0 100.0 91.7 91.7 91.7 91.7 Detected Category 2 (%) 84.2 84.2 81.6 79.0 79.0 60.5 Detected Volcanoes (%) 79.0 77.7 75.1 73.4 70.4 63.1 False Alarms per image 42.8 35.5 28.7 21.4 13.2 5.0 False Alarms per (d) HET38 Operating point: OP1 OP2 OP3 OP4 OP5 OP6 Threshold: 0.75 0.80 0.85 0.90 0.95 0.99 Detected Category 1 (%) 90.3 87.1 87.1 83.9 79.0 64.5 Detected Category 2 (%) 84.6 82.4 80.2 78.7 75.0 62.5 Detected Volcanoes (%) 74.1 72.0 69.8 66.2 60.9 47.8 False Alarms per image 50.2 43.7 37.1 29.7 20.5 10.6 False Alarms per 10 4 89.2 77.6 65.9 52.7 36.4 18.8 Figure 10. Performance of the baseline system at various operating points along the FROC curve. (a) HOM4 (b) HET5 Alarms km Detection Rate baseline - tts baseline - loi 1-class nn - lox 1-class nn - loi 2-class nn - lox 2-class nn - loi Alarms Detection Rate baseline - tts baseline - loi 1-class nn - lox 1-class nn - loi 2-class nn - lox 2-class nn - loi baseline - tts baseline - loi 1-class nn - lox 1-class nn - loi 2-class nn - lox 2-class nn - loi Figure 11. Performance results of svd-gauss, 1-class nearest neighbor, and 2-class nearest-neighbor algorithms under the leave-out-image (LOI) and leave-out-example (LOX) training paradigms. (a) Results on a set of four homogeneous images from one area of the planet. (b) Results on a set of five heterogeneous images selected from different areas of the planet. Refer to the text for an interpretation of the results. test example is assigned to the same class as its nearest neighbor in the reference library. The other was a one-class version in which the reference library contains only positive examples (volcanoes); an unknown example is assigned to the volcano class if it is similar enough to some member of the reference library. Performance was evaluated on both the HOM4 and HET5 datasets using the three classifiers (baseline, 1-class nearest neighbor, and 2-class nearest neighbor) and two training paradigms (LOI and LOX). The results are shown in Figure 11. For computational reasons, the baseline method was trained and tested on the same data (TTS) rather than leaving out an example. The effect of including the test example in the training set is minimal in this case since one example has little effect on the class-conditional mean and covariance estimates. The following is the key observation: on HET5 the baseline and 2-class nearest-neighbor algorithms work significantly better under the LOX training paradigm than under the LOI training paradigm; however, the 1-class algorithm works the same under both training paradigms. If having knowledge about the local volcanoes were the critical factor, the 1-class algorithm should have worked significantly better under LOX than under LOI. Instead we conclude that access to the local non- volcanoes is the critical factor. The 1-class algorithm completely ignores the non- volcanoes and hence does not show any difference between LOI and LOX. The other methods do use the non-volcanoes, and these show a dramatic improvement under LOX. On HOM4 there is little difference between the LOI and LOX results. Since these images are from the same area of the planet, the appearance of the non-volcanoes is similar from image to image. Thus, leaving out one example or leaving out one image from the training set does not have much effect. The non-volcanoes in HET5 and other heterogeneous image sets vary considerably from image to image and this may be the source of the degradation in performance (the training data is inadequate for learning the non-volcano distribution). 4. Project Status Participating in the development of the JARtool system were two planetary geologists (Aubele and Crumpler) who were members of the Volcanism Working Group of the Magellan Science team and principal investigators in NASA's Planetary Mapping and Venus Data Analysis Programs. The geologists have been evaluating the JARtool approach both in terms of the scientific content provided by the analyzed images and as a tool to aid in further cataloging. From the planetary geologists' point of view, the primary goal was to achieve annotation of 80% or more of the small volcanoes in the analyzed datasets. A secondary goal was to obtain accurate diameter estimates for each volcano. Locating different morphologic types of small volcanoes was also of interest. However, it was recognized up front that some of the types would be easy to detect and some would be difficult (both for human experts and for algorithms). To the geologists, the system should be considered a success if it detects a high percentage of the "easy" volcanoes (category 1 and 2). Our test results indicate that this level of performance is achieved on homogeneous image sets. However, we have not succeeded in developing a reliable method for measuring volcano diameters. Hence, sizing capability is not included in the fielded system. Our experiments and those of the scientists have indicated that the choice of operating point will vary across different areas of the planet, dependent on factors such as terrain type and local volcano distributions. Hence, the operating point is left "open" for the scientists to choose. Although the original intent was for the JARtool system to provide a fully-automated cataloging tool, it appears that the system will be most useful as an "intelligent assistant" that is used in an interactive manner by the geologists. The capabilities of the system were recently expanded through integration of the Postgres database (Stonebraker and Kemnitz, 1991). A custom query tool supports arbitrary SQL queries as well as a set of common "pre-canned" queries. JARtool is also being evaluated for use in other problem domains. Researchers or scientists who are interested in the software can direct inquiries to [email protected]. 5. Lessons Learned and Future Directions Real-world applications of machine learning tend to expose practical issues which otherwise would go unnoticed. We share here a list of "lessons learned" which can be viewed as a "signpost" of potential dangers to machine learning practitioners. In addition, for each "lesson learned" we discuss briefly related research opportunities in machine learning and, thus, provide input on what topics in machine learning re-search are most likely to have practical impact for these types of large-scale projects in the future. 1. Training and testing classifiers is often only a very small fraction of the project effort. On this project, certainly less than 20%, perhaps as little as 10% effort was spent on classification aspects of the problem. This phenomenon has been documented before across a variety of applications (Langley and Simon, 1994; Brodley and Smyth, 1997). Yet, this directly contradicts the level of effort spent on classification algorithms in machine learning research, which has traditionally focused heavily on the development of classification algorithms. One implication is that classification technology is relatively mature and it is time for machine learning researchers to address the "larger picture." A difficulty with this scenario is that these "big picture" issues (some of which are discussed below) can be awkward to formalize. 2. A large fraction of the project effort (certainly at least 30%) was spent on "fea- ture engineering," i.e., trying to find an effective representation space in which to apply the classification algorithms. This is a very common problem in applications involving sensor-level data, such as images and time-series. Unfortunately, there are few principled design guidelines for "feature engineering" leading to much trial-and-error in the design process. Commonly used approaches in machine learning and pattern recognition are linear projection techniques (such as PCA) and feature selection methods. Non-linear projection techniques can be useful but are typically computationally complex. A significant general problem is the branching factor in the search space for possible feature representations. There are numerous open problems and opportunities in the development of novel methods and systematic algorithms for feature extraction. In particular, there is a need for robust feature extraction techniques which can span non-standard data types, including mixed discrete and real-valued data, time series and sequence data, and spatial data. 3. Real-world classification systems tend to be composed of multiple components, each with their own parameters, making overall system optimization difficult if not impossible given finite training sets. For JARtool, there were parameters associated with FOA, feature extraction, and classification. Joint optimization of all of these parameters was practically impossible. As a result many parameters (such as the window size for focus of attention) were set based on univariate sensitivity tests (varying one parameter while keeping all others fixed at reason-able values). Closer coupling of machine learning algorithms and optimization methods would appear to have significant potential payoffs in this context. 4. In many applications classification labels are often supplied by experts and may be much noisier than initially expected. At the start of the volcano project, we believed the geologists would simply tell us where all the volcanoes were in the training images. Once we framed the problem in an ROC context and realized that the resolution of the images and other factors led to inherent ambiguity in volcano identification, we began to understand the noisy, subjective nature of the labeling process. In fact, the geologists were also given cause to revise their opinions on the reliability of published catalogs. As real-world data sets continue to grow in size, one can anticipate that the fraction of data which is accurately labeled will shrink dramatically (this is already true for many large text, speech, and image databases). Research areas such as coupling unsupervised learning with supervised learning, cognitive models for subjective labeling, and active learning to optimally select which examples to label next, would appear to be ripe for large-scale application. 5. In applying learning algorithms to image analysis problems, spatial context is an important factor and can considerably complicate algorithm development and evaluation. For example, in testing our system we gradually realized that there were large-scale spatial effects on volcano appearance, and that training a model on one part of the planet could lead to poor performance elsewhere. Conversely, evaluating model performance on images which are spatially close can lead to over-optimistic estimates of system performance. These issues might seem trivial in a machine learning context where independence of training and test sets is a common mantra, yet the problem is subtle in a spatial context (How far away does one have to go spatially to get independent data?) and widely ignored in published evaluations of systems in the image analysis and computer vision communities. Thus, there is a need to generalize techniques such as cross-validation, bootstrap, and test-set evaluation, to data sources which exhibit dependencies (such as images and sequences). A common theme emerging from the above "lessons" is that there is a need for a systems viewpoint towards large-scale learning applications. For example, in ret- rospect, it would have been extremely useful to have had an integrated software infrastructure to support data labeling and annotation, design and reporting of experiments, visualization, classification algorithm application, and database support for image retrieval. (For JARtool development, most of these functions were carried out within relatively independent software environments such as standalone C programs, Unix shell scripts, MATLAB, SAOimage, and so forth). Development of such an integrated infrastructure would have taken far more resources than were available for this project, yet it is very clear that such an integrated system to support application development would have enabled a much more rapid development of JARtool. More generally, with the advent of "massive" data sets across a variety of disci- plines, it behooves machine learning researchers to pay close attention to overall systems issues. How are the data stored, accessed, and annotated? Can one develop general-purpose techniques for extracting feature representations from "low-level" data? How can one best harness the prior knowledge of domain experts? How can success be defined and quantified in a way which matches the user's needs? 6. Conclusion The Magellan image database is a prime example of the type of dataset that motivates the application of machine learning techniques to real-world problems. The absence of labeled training data and predefined features imposes a significant challenge to "off the shelf" machine learning algorithms. JARtool is a learning-based system which was developed to aid geologists in cataloging the estimated one million small volcanoes in this dataset. The system is trained for the specific volcano-detection task through examples provided by the geologists. Experimental results show that the system approaches human performance on homogeneous image sets but performs relatively poorly on heterogeneous sets in which images are selected randomly from different areas of the planet. The effect on system performance of a particular classification algorithm was found to be of secondary importance compared to the feature extraction problem. Acknowledgments The research described in this article has been carried out in part by the Jet Propulsion California Institute of Technology, under contract with the National Aeronautics and Space Administration. Support was provided by the NASA Office of Advanced Concepts and Technology (OACT - Code CT), a JPL DDF award, NSF research initiation grant IRI 9211651, and a grant from the Swedish Foundation for International Cooperation in Research and Higher Education (Lars Asker). We would like to thank Michael Turmon for his help and for performing some of the experiments. We would also like to thank Saleem Mukhtar, Maureen Burl, and Joe Roden for their help in developing the software and user-interfaces. The JARtool graphical user interface is built on top of the SAOtng image analysis package developed at the Smithsonian Astrophysical Society (Mendel et al., 1997). Notes 1. The nominal pixel spacing in the highest resolution Magellan data products is 75m, but this image was resized slightly. 2. One difference is that the area under an FROC curve cannot be interpreted in the same way as for a true ROC curve. --R Small domes on Venus: characteristics and origins. Personal communication. Applying classification algorithms in practice. Trainable cataloging for digital image libraries with applications to volcano detection. Personal communication. Human visual perception and ROC methodology in medical imaging. Experts in Uncertainty. Detection of circular geological features using the Hough transform. Volcanoes and centers of volcanism on Venus. Pattern Classification and Scene Analysis. A learning approach to object recognition: applications in science image analysis. Query by image and video content - the QBIC system Introduction to Statistical Pattern Recognition. Signal Detection Theory and Psychophysics. Small volcanic edifices and volcanism in the plains of Venus. Machine learning for the detection of oil spills in satellite radar images. Applications of machine learning and rule induction. Signal Detection Theory: A User's Guide. SAOimage: the next generation. Maximum likelihood detection of faces and hands. Magellan: Radar Performance and Data Products. Introduction to the special section on digital li- braries: representation and retrieval Behavioral Decision Theory: A New Approach. Analysis and visualization of classifier performance: comparison under imprecise class and cost distributions. Remote Sensing for Digital Image Analysis. Magellan Mission Summary. Efficient pattern recognition using a new transformation distance. Low dimensional procedure for the characterization of human faces. The Hough transform applied to SAR images for thin line detection. Signal detection theory: valuable tools for evaluating inductive learning. Global distribution and characteristics of coronae and related features on Venus - Implications for origin and relation to mantle processes The POSTGRES next generation database- management system Image feature reduction through spoiling: its application to multiple matched filters for focus of attention. Using discriminant eigenfeatures for image retrieval. The design of multistage separable planar filters. Eigenfaces for recognition. Personal communication. Recognition of volcanoes using correlation methods. --TR The Hough transform applied to SAR images for thin line detection Introduction to statistical pattern recognition (2nd ed.) Signal detection theory: valuable tools for evaluating inductive learning The POSTGRES next generation database management system Applications of machine learning and rule induction Introduction to the Special Section on Digital Libraries Using Discriminant Eigenfeatures for Image Retrieval Photobook Machine Learning for the Detection of Oil Spills in Satellite Radar Images Remote Sensing Applying classification algorithms in practice Query by Image and Video Content Efficient Pattern Recognition Using a New Transformation Distance --CTR Steve Chien , Rob Sherwood , Daniel Tran , Benjamin Cichy , Gregg Rabideau , Rebecca Castano , Ashley Davies , Rachel Lee , Dan Mandl , Stuart Frye , Bruce Trout , Jerry Hengemihle , Jeff D'Agostino , Seth Shulman , Stephen Ungar , Thomas Brakke , Darrell Boyer , Jim Van Gaasbeck , Ronald Greeley , Thomas Doggett , Victor Baker , James Dohm , Felipe Ip, The EO-1 Autonomous Science Agent, Proceedings of the Third International Joint Conference on Autonomous Agents and Multiagent Systems, p.420-427, July 19-23, 2004, New York, New York Rie Honda , Shuai Wang , Tokio Kikuchi , Osamu Konishi, Mining of Moving Objects from Time-Series Images and its Application to Satellite Weather Imagery, Journal of Intelligent Information Systems, v.19 n.1, p.79-93, July 2002 Steve Chien , Rob Sherwood , Daniel Tran , Benjamin Cichy , Gregg Rabideau , Rebecca Castao , Ashley Davies , Dan Mandl , Stuart Frye , Bruce Trout , Jeff D'Agostino , Seth Shulman , Darrell Boyer , Sandra Hayden , Adam Sweet , Scott Christa, Lessons learned from autonomous sciencecraft experiment, Proceedings of the fourth international joint conference on Autonomous agents and multiagent systems, July 25-29, 2005, The Netherlands Jianming Liang , Marcos Salganicoff, On the medical frontier: the 2006 KDD Cup competition and results, ACM SIGKDD Explorations Newsletter, v.8 n.2, p.39-46, December 2006 Steve Chien , Rob Sherwood , Gregg Rabideau , Rebecca Castano , Ashley Davies , Michael Burl , Russell Knight , Tim Stough , Joe Roden , Paul Zetocha , Ross Wainwright , Pete Klupar , Jim Van Gaasbeck , Pat Cappelaere , Dean Oswald, The Techsat-21 autonomous space science agent, Proceedings of the first international joint conference on Autonomous agents and multiagent systems: part 2, July 15-19, 2002, Bologna, Italy George Gigli , loi Boss , George A. Lampropoulos, An optimized architecture for classification combining data fusion and data-mining, Information Fusion, v.8 n.4, p.366-378, October, 2007 Tom Fawcett , Foster Provost, Activity monitoring: noticing interesting changes in behavior, Proceedings of the fifth ACM SIGKDD international conference on Knowledge discovery and data mining, p.53-62, August 15-18, 1999, San Diego, California, United States Chandrika Kamath , Erick Cant-Paz , Imola K. Fodor , Nu Ai Tang, Classifying of Bent-Double Galaxies, Computing in Science and Engineering, v.4 n.4, p.52-60, July 2002 Foster Provost , Ron Kohavi, Guest Editors Introduction: On Applied Research in MachineLearning, Machine Learning, v.30 n.2-3, p.127-132, Feb./ March, 1998 Miroslav Kubat , Robert C. Holte , Stan Matwin, Machine Learning for the Detection of Oil Spills in Satellite Radar Images, Machine Learning, v.30 n.2-3, p.195-215, Feb./ March, 1998 Paul Stolorz , Peter Cheeseman, Onboard Science Data Analysis: Applying Data Mining to Science-Directed Autonomy, IEEE Intelligent Systems, v.13 n.5, p.62-68, September 1998 Dennis Decoste , Bernhard Schlkopf, Training Invariant Support Vector Machines, Machine Learning, v.46 n.1-3, p.161-190, 2002 M. A. Maloof , P. Langley , T. O. Binford , R. Nevatia , S. Sage, Improved Rooftop Detection in Aerial Images with Machine Learning, Machine Learning, v.53 n.1-2, p.157-191, October-November

automatic cataloging;data mining;detection of natural objects;principal components analysis;JARtool;volcanoes;trainable;learning from examples;pattern recognition;venus;machine learning;large image databases;Magellan SAR

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Interference-Minimizing Colorings of Regular Graphs.

Communications problems that involve frequency interference, such as the channel assignment problem in the design of cellular telephone networks, can be cast as graph coloring problems in which the frequencies (colors) assigned to an edge's vertices interfere if they are too similar. The paper considers situations modeled by vertex-coloring d-regular graphs with n vertices using a color set 1, 2,..., n, where colors i and j are said to interfere if their circular distance $\min \{ | i-j | , n- | i-j | \}$ does not exceed a given threshold value $\alpha$. Given a d-regular graph G and threshold $\alpha$, an interference-minimizing coloring is a coloring of vertices that minimizes the number of edges that interfere. Let $I_\alpha (G)$ denote the minimum number of interfering edges in such a coloring of $G$. For most triples $(n, \alpha ,d),$ we determine the minimum value of $I_\alpha (G)$ over all d-regular graphs and find graphs that attain it. In determining when this minimum value is 0, we prove that for $r \geq 3$ there exists a d-regular graph G on n vertices that is r-colorable whenever $d \leq (1- \frac{1}{r}) n-1$ and nd is even. We also study the maximum value of $I_\alpha (G)$ over all d-regular graphs and find graphs that attain this maximum in many cases.

Introduction This paper is motivated by telecommunication problems such as the design of planar regions for cellular telephone networks and the assignment of allowable frequencies to the regions. In our graph abstraction, vertices are regions, edges are pairs of contiguous regions, and colors correspond to frequencies. We presume that every region has the same number d of neighbors, which leads to considering degree-regular graphs. Interference occurs between two regions if they are neighbors and their frequencies lie within an interference threshold. We adopt the simplifying assumption that the number of colors available equals the number n of regions, and let ff denote the threshold parameter so that colors i and j in f1; ng interfere precisely when their circularly-measured scalar distance is less than or equal to ff. Precedents for the use of circularly-measured distance in graph coloring include Vince (1988) and Guichard and Krussel (1992). Our formulation leads to several interesting graph-theoretic problems. One is to determine for any given d-regular graph G and threshold ff the minimum number I ff (G) of interfering edges over the possible colorings of G. Another is: given parameters n; ff, and d, determine the minimum and maximum values of I ff (G) and find graphs G that attain these values. We focus on the latter problem. More specifically, let G(n; d) denote the set of undirected d-regular graphs on n vertices, which have no loops or multiple edges, but may be disconnected. We wish to determine the (global) minimum interference level '(n; ff; d), which is the minimum of I ff (G) over G(n; d). For comparison purposes, we also wish to determine the (global) minimax interference level L(n; ff; d), which is the maximum of I ff (G) over G(n; d). This latter problem measures how badly off you would be if an adversary gets to choose G 2 G(n; d), and you can then color G to minimize interference. Our graph-theoretic model is an approximation to the frequency assignment problem for cellular networks studied in Benveniste et al. (1995). In that paper the network of cellular nodes is viewed as vertices of a hexagonal lattice in R 2 , and the graph G is specified by a choice of sublattice 0 of , with being the index of the sublattice 0 in . More precisely, the vertices of G are cosets of and we draw an edge between two cosets if the cosets are "close" in the sense that they contain vectors v, v 0 respectively with where jj \Delta jj is a given norm on R 2 and x is a cutoff value. Such graphs 1 G are d-regular for 1 The graph G represents a fundamental domain of . In the cellular terminology a fundamental domain for called a "reuse group." More generally a "reuse group" is a collection of contiguous cells that exhausts all frequencies, with no two cells in the group using the same frequency. some value of d; the usual nearest-neighbors case gives et al. (1995). The frequency spectrum is also divided into cosets (modulo n), and nodes in the same coset (mod 0 ) are assigned a fixed coset of frequencies (mod n). In cellular problems the graph G is fixed (depending on 0 ). Typical parameters under consideration are and n=ff about 2 or 3. From this standpoint the quantities '(n; ff; d) and L(n; ff; d) represent lower and upper bounds for attainable levels of interference. Related coloring problems motivated by the channel assignment problem are studied in Hale (1980), Cozzens and Roberts (1982), Bonias (1991), Liu (1991), Tesman (1993), Griggs and Liu (1994), Raychaudhuri (1994), Troxell (1996) and Guichard (1996) among others. Roberts (1991) surveys the earlier part of this work. Factors that distinguish prior work from the present investigation include our focus on regular graphs and the inevitability of interference when certain relationships hold among n; ff and d. Our main results give near-optimal bounds for '(n; ff; d) and L(n; ff; d) and identify d-regular graphs and colorings that attain extremal values. Many interference-minimizing designs use only a fraction of the available colors or frequencies. The most common number of colors used in these optimal designs is which is the maximum number of mutually noninterfering colors from f1; ng at threshold ff. Detailed statements of theorems for '(n; ff; d) and L(n; ff; d) appear in Section 2. Proofs follow in Sections 3 to 7. In the course of our analysis we derive a graph-theoretic result of interest in its own right, which is a condition for the existence of a d-regular graph having chromatic number - r. Theorem 1.1. If r - 3, then G(n; d) contains an r-colorable graph if nd is even and r This result is proved in Section 5, and the proof can be read independently of the rest of the paper. Note that if nd is odd then G(n; d) is the empty set. We preface the results in the next section with a few comments to indicate where we are headed. The case corresponds to no interference because the number of available colors equals the number of vertices, and therefore '(n; 0; d) = L(n; 0; d) = 0. We assume that ff - 1 in the rest of the paper. For degrees near 0 or n, namely 1, the set G(n; d) contains only one unlabelled graph, so these cases are essentially trivial. We note at the end of Section 4 that Our first main result in the next section, Theorem 2.1, applies to degree 2 and shows that most values of ' and L for exception is that L(n; 2; 2) is approximately n=3. Subsequent results focus on d - 3, where we use the maximum number of noninterfering colors fl to express the results. The case because then all colors interfere with each other, so that #(edges of d) for most values of (n; ff; d) is approximately nd Moreover, L(n; ff; d) = 0 whenever fl ? d, whereas if n is much larger than d, and d is somewhat larger than fl, then L(n; ff; d) is approximately nd=(2fl). Extremal graphs which attain '(n; ff; d) when ' ? 0 are usually connected, and the associated coloring can often be achieved using fl noninterfering colors. On the other hand, graphs that attain L(n; ff; d) when L ? 0 are usually disconnected and contain many copies of the complete graph K d+1 . There are exceptions, however. Our results imply that there is often a sizable gap between the values of ' and L. The smallest instance of l ! L occurs at (n; ff; d) 2. Figure 1.1 shows the two graphs in G(6; 2) with interference-minimizing colorings for 2. Figure 1.1 about here A qualitative comparison of the regions where ' and L equal 0 and are positive is given in Figure 1.2, where the coordinates are d=n and fl=n. Figure 1.2 about here 2. Main Results An undirected graph is simple if it has no loops or multiple edges. Let G(n; d) denote the set of d-regular graphs on n vertices which are simple but which are not necessarily connected. ng be a set of n colors with circular distance measure and let ff 2 f0; be the threshold-of-interference parameter. A coloring of the vertex set V (G) of graph d) is a [n]. The interference I ff (G; f) of coloring f of G at threshold ff is I ff (G; f) := jffx; yg The minimum interference in G at threshold ff is I ff (G) := min f :V (G)![n] I ff (G; f) : We study the (global) minimum interference level '(n; ff; d) := min G2G(n;d) I ff (G) (2.1) and the (global) minimax interference level G2G(n;d) I We first note restrictions on the parameter space. Since all graphs in G(n; d) have ndedges, it follows that n and d cannot both be odd : (2.3) We restrict attention to the threshold range because ff - nimplies that all colors interfere. Thus Our first result concerns ' and L for degree 2. Theorem 2.1. Let 2. (a) For all (b) For all fl - 3, (c) If This is proved in Section 3. We now consider d in the range for the minimum interference level '. The cases of are treated separately. We obtain an almost complete answer for 2. Theorem 2.2. Suppose that 2. (a) If n is even, then and nis even, or nand d are both odd and d is even : (b) If n is odd, then and (c) If n is odd and in the remaining range n\Gamma2ff - d - n, then '(n; ff; d) - d. Furthermore: there is an integer 2s (ii) '(n; ff; d) - d 2 is even, and there is an integer 4s Case (c) above is the only case not completely settled. Instances of it are illustrated in Figure 2.1. The number beside each vertex clump gives the color assigned to those vertices, and the number on a line between noninterfering clumps is the number of edges between them. Case analyses, omitted here, show that no improvements are possible in part (c) of the theorem when Given n - 21, (i) has three realizations, namely '(15; 5; 0, (ii) has only the realization at the bottom of Figure 2.1, and cases. Figure 2.1 about here We remark that the bounds on '(n; ff; d) for d ? n are obtained using a variant of Turan's theorem on extremal graphs (Tur'an, 1941; Bondy and Murty, 1976, p. 110). Theorem 2.2 is proved in Section 4. We now consider the minimum interference level ' when fl - 3. To handle this case we use Theorem 1.1, which is proved in Section 5. Let p and q be the unique nonnegative integers that satisfy that is c and Our bounds for are given in the next two theorems for respectively, and are proved in Section 6. The case is somewhat simpler. Theorem 2.3. Suppose that fl - 3 and that fl divides n, i.e. (a) If d - (b) If d ? is odd or if even and p is even even and p is odd : Theorem 2.4 Suppose that fl - 3 and fl doesn't divide n, i.e. q - 1. (a) If d (b) If d - where even and p is even even and p is odd : We turn next to results for the minimax interference level L. We first distinguish cases Theorem 2.5. Suppose that 3 - d - 2. Then: (a) L(n; ff; d) (b) L(n; ff; d) ? 0 for The only cases in the parameter range 1 - ff - n\Gamma 1 and fl - 2 not settled by this theorem are those with Both occur in this exceptional case, e.g. for Our final main result provides bounds for L. Set c and In view of Theorem 2.5 we consider only the range that 2 - fl - d. Theorem 2.6. Suppose that 3 - d - and In the special case that d can be written more simply as nd This applies in particular when which case the upper and lower bounds coincide, yielding (1.1). If n is substantially larger than d, and d is somewhat larger than fl, then L is closely approximated by nd Theorems 2.5 and 2.6 are proved in Section 7. 3. Elementary Facts: Theorem 2.1 We derive general conditions that guarantee graphs (Theorem 2.1). Lemma 3.1. If 1 - ff ! n (a) '(n; ff; d) and (b) '(n; ff; d) 2 and n is even: (3.2) Proof. (a) Given d ng and consider the coloring for every i. We construct a suitable G starting with the edge set If n is odd, or if n is even and d is odd, let every vertex has degree d and every edge has D ? ff, so '(n; ff; d) = 0. If n and d are both even, so ff - (n ng . Again, every vertex has degree d and every edge has D ? ff, so '(n; ff; d) = 0. (b) Let -G denote the chromatic number of the graph G. The definition implies that: If n is even and d - n d) contains a bipartite graph with n=2 vertices in each part, so (b) follows from (3.3), since fl - 2. \Xi We remark that the construction in part (a) uses all n colors, and when d same construction gives many interfering edges. It is natural to consider the opposite extreme, which is to use only a maximal set of noninterfering colors. This leads to part (b). The restriction in part (b) that n be even is crucial, because no d-regular bipartite graph exists for odd n. Indeed there are exceptions where '(n; ff; d) ? 0 for some d ! n 2 with n odd: see Theorems 2.1 and 2.2. These exceptions occur when but are not an issue for fl - 3. We obtain bounds on the minimax interference level L using the following well-known bound for the chromatic number -G of a graph G. Proposition 3.1. For every finite simple graph G, where \Delta G is the maximum degree of a vertex of G. Furthermore -G - \Delta G provided that no connected component of G is an odd cycle or a complete graph. Proof. Brooks (1941); Bondy and Murty (1976, pp. 118 and 122). \Xi This result immediately yields the following condition for the minimax interference level Lemma 3.2. If 1 - ff ! n Proof. The definition of L(n; ff; d) gives for a d-regular graph, (3.5) follows from Proposition 3.1 via (3.6). \Xi Proof of Theorem 2.1. (a) Since follows from (3.2) if n is even, and from if n is odd and ff - n (b) follows from Lemma 3.2. (c) Given every graph in G(n; 2) is a sum of vertex-disjoint cycles. Suppose 1. Then an even cycle has minimum interference 0, a 3-cycle has minimum interference 1, and an odd cycle with five or more vertices has minimum interference 0 or 1. It follows that one 4-cycle), and L 2 2. The last case uses M \Gamma 1 3-cycles and one 5-cycle. When the 5-cycle's vertices are colored successively as 1, ff ff, it has no interference if [n so 1 in this case. More generally, suppose one vertex of the 5-cycle is colored 1. Its neighbors must have colors in [ff to avoid interference. Then their uncolored neighbors, which are adjacent, must have colors in [2ff to avoid interference. This set has which is - ff if (2n \Gamma 4)=5 - ff. Hence 4. Minimal Interference Level: Theorem 2.2 We prove Theorem 2.2 in this section. The ranges stated where '(n; ff; d) = 0 follow from Lemma 3.1, so the main content of parts (a) and (b) of Theorem 2.2 concerns the values '(n; ff; d) for d ? n 2 . To obtain these we use a variant of Tur'an's theorem (Tur'an, 1941; Bondy and Murty, 1976, p. 110), which we state as a lemma. An application of the lemma at the end of the section yields the exact value of L(n; ff; well as '(n; ff; 1). Recall that an equi-t-partition of a vertex set V is a partition fV Lemma 4.1. The maximum number of noninterfering edges in the complete graph K n with vertex set V and threshold parameter ff is attained only by a coloring [n] that has are in different parts of an equi-fl-partition of V . Proof. Suppose that a coloring f of the complete graph K n has f i vertices of color i ab denote the number of vertices of colors other than a and b that interfere with a and not b. If all color-i vertices are recolored j, the net increase in interference is f i (m vertices are recolored i, the net increase in interference if f Hence at least one of the recolorings does not increase interference. Continuing this recoloring process implies that noninterference in K n is maximized by a fl-partite partition of V such that D(f(x); f(y)) ? ff whenever x and y are in different parts of the partition. Tur'an's theorem then implies that maximum noninterference obtains only when the partition is an equi-fl-partition. \Xi We can assume without loss of generality that the coloring f found in Lemma 4.1 is constant on each part of an equi-fl-partition, with f(V flg. If interfering edges are then dropped from K n , we obtain a complete equi-fl-partite graph with zero interference and chromatic number fl. This graph is regular if and only if fl divides n and each part of the partition has n=fl vertices. Proof of Theorem 2.2. Throughout this proof We consider first (a) and (b). The ranges given where '(n; ff; d) = 0 come from Lemma 3.1. So assume now that d ? n . Let G 0 be a complete bipartite graph fA; Bg such that Lemma 4.1 implies that two-coloring G 0 using noninterfering colors for A and B uniquely maximizes the number of edges with no interference when 2. Therefore ' - nd=2 \Gamma jAjjBj . (a) Suppose n is even. If n=2 and d are odd, the number of edges needed within each part of G 0 to increase all degrees to d is (n=2)(d \Gamma n=2)=2, which is an integer since d \Gamma n=2 is even. It follows that if n=2 is even, or if n=2 and d are odd, then If instead n=2 is odd, and d is even, then (n=2)(d \Gamma n=2) is odd, G 0 is not part of any graph in G(n; d), and ' ? replacing G 0 with a complete bipartite graph G 1 with bipartition fA Beginning with G 1 , each vertex in A 0 requires more degrees to have degree d, and each vertex in B 0 requires edges added to have degree d. Both d and are even, so edge additions as needed can be made within A 0 and B 0 to obtain 1 in this case; and (2.9) is proved. (b) Suppose n is odd, so d is even by (2.3). Beginning with G 0 , each of the (n vertices in A requires d \Gamma (n \Gamma 1)=2 more incident edges added to have degree d, and each of the incident edges added to have degree d. Each of f(n contains an even integer, so we can make the required additions of edges within A and B. Hence It remains to prove (c), which has three parts (i)-(iii). Assume henceforth that n is odd even because n is odd. Augmented equi-bipartite graphs, illustrated at the top of Figure 2.1, show that ' - d=2 since they require d=2 edges within the 1)=2-vertex part to obtain degree d for every vertex. Sometimes case in point is the largest possible ff for Suppose 1g. Each vertex in the color set [n] has exactly two others for which D ? ff, and the graph of noninterfering colors is an n-cycle whose successive colors are 1)=2. If every color were assigned to some vertex in G 2 G(n; d), there would be at least n(d \Gamma 2)=2 interference edges. But n(d must avoid at least one color to attain '. Deletion of one color from the n-cycle of noninterfering colors breaks the cycle and leaves the noninterference graph Because all x i colors interfere with each other, and all y i colors interfere with each other, we can presume that f uses only one x i and an adjacent y j . This yields the augmented bipartite structure of the preceding paragraph, and it follows from maximization of between-parts edges that This completes the proof of (iii). For (i) and (ii), assume ff ! (n \Gamma 3)=2 and consider an odd r - 5 sequence of colors c 1 , 1. The tightest such sequence has to color k. It follows that the final color c r can be chosen not to interfere with c i.e., if We usually consider the smallest such odd r - 5 because this allows the the largest d values. Our approach, illustrated on the lower part of Figure 2.1, is to assign clumps of vertices to the c i in such a way that all edges for G 2 G(n; d) are between adjacent clumps on the noninterference color cycle c Suppose (4:2) holds for a fixed odd r - 5. We assume that r ! n because the ensuing analysis requires this for d - 3. Let a and b be nonnegative integers that satisfy We prove (i), then conclude with (ii). The analysis for (i) splits into three cases depending on the parity of a and br=4c. Case 1: a odd Case 2: a even, br=4c odd Case 3: a even, br=4c even. Because n is odd, Case 1 requires b to be even and Cases 2 and 3 require b to be odd. Case 1. Given an odd a, we partition the n vertices into b clumps of a vertices each clumps of a vertices each. The clumps are assigned to colors in the noninterference cycle so that the clumps of each type are contiguous. Cases for are illustrated at the top of Figure 4.1. We begin at the central (top) a clump and proceed symmetrically in both directions around the color cycle, assigning between-clumps edges as we go so that all vertices end up with degree 2a. The required edges into the next clump encountered are distributed as equally as possible to the vertices in that clump. When we get into the clumps with a+ 1 vertices, the number of between-clumps edges needed will generally be less than the maximum possible number of (a . Numbers of between-clumps edges used to get degree 2a for every vertex are shown on the noninterference lines between the c i on Figure 4.1. Figure 4.1 about here The preceding construction yields '(n; ff; d) = 0 for even d is less than 2a, say we modify the procedure by using fewer between-clumps edges for the required vertex degrees: clump sizes are unchanged. Because n=r yields the contradiction that n ! ra, it follows for Case 1 that Case 2. With a even and br=4c odd, we have b odd and r 2 f5; 7; 13; 15; :g. In this case we assign a \Gamma 1 vertices to c 1 and proceed in each direction around the c i cycle, assigning a, a vertices to the next (r \Gamma 1)=2 c i in order. The penultimate number a 0 equals a if r 2 f5; 13; and is a :g. The ultimate number a in chosen so that there are [n \Gamma (a \Gamma 1)]=2 vertices (excluding the a \Gamma 1 for 1 ) on each side of the color cycle. If a a 1)=2. The two cases are shown on the lower left of Figure 4.1 with numbers of between-clumps edges that give degree d = 2a for every vertex. If even d is less than 2a, fewer edges are used, as needed, down the two sides. As in Case 1, we get Case 3. With a even and br=4c even, we have b odd and r 2 f9; 11; 17; 19; 25; Here we assign a+ 1 vertices to c 1 and proceed with a, a \Gamma 1, a, a+ 1, a, a \Gamma 1, a, a+ a vertices assigned to the next (r \Gamma 1)=2 c i in each direction away from c 1 . We get a and a :g. The two cases are shown on the lower right of Figure 4.1. As before, if d - (2=r)n. This completes the proof of (i), after defining s by 1. We have also checked that the construction used here cannot yield unless the conditions of (i) hold. There is however one other set of circumstances where this construction yields a value of r n, and these circumstances are exactly the hypotheses of (ii), namely:? d - 8; d=2 is even is odd since d=2 is even) In this case we partition the vertices into (r +1)=2 clumps of vertices each and (r \Gamma 1)=2 clumps of d=2 The clumps are arranged around the noninterference color cycle c 1 , c We use all possible between-clumps edges. This gives degree d for every vertex except those in the c 2 clump, which has incoming edges from c 1 and c 3 . The degree total for c 2 should be d, so we need to add to get degree d for each c 2 vertex. Prior to the additions, each c 2 vertex has degree d \Gamma 2 by our equalization construction, so the additions can be made by a complete cycle within the clump. It follows that ' - d=2 \Gamma 1, proving (ii). \Xi We conclude this section by noting that the modified Tur'an's theorem (Lemma 4.1) easily allows us to completely settle the case of degree Corollary 4.2. For c Proof. Write so An equi-fl-partition of an n vertex set has? ! q parts, each with each with p vertices. Now the unique graph G 2 G(n; applying Lemma 4.1, we have p! which is (4.4). \Xi 5. Chromatic Number Bound: Theorem 1.1 This section gives a self-contained proof of Theorem 1.1. We first recall two preliminary facts, stated as propositions. Proposition 5.1. (Dirac (1942)) Let G be a simple graph. If every vertex of G is of degree at least jV (G)j=2, then G is Hamiltonian, that is, G has a cycle of length jV (G)j. Proof. See Bondy and Murty (1976), p. 54. \Xi Recall that a matching in a simple graph G is a subset of mutually vertex-disjoint edges of G. A matching is perfect if every vertex in G is on some edge of the matching. The following is a consequence of a well-known theorem of Hall (1935). Proposition 5.2. (Marriage Theorem) If G is a d-regular bipartite graph with d ? 0, then G has a perfect matching. Proof. See Bondy and Murty (1976), p. 73. \Xi We study the function OE(n; d; r) defined by 1 if there exists an n-vertex d-regular r-colorable graph, When OE(n; d; denote such a d-regular r-colorable (that is, r-partite) graph having n vertices. We consider only values in which nd is even. Our first observation is that because an r-colorable graph is also (r + 1)-colorable, The purpose of the next two lemmas is to prove that OE(n; d; r) is monotone when r - 3 is held fixed and d varies over values where nd is even. Lemma 5.1 (a) If d - nand if either r - 3 or (b) If d - n If in addition n is even, then Proof. (a) Suppose that n is even. The inequality (5.1) implies that it is enough to show OE(n; d; We use reverse induction on d - n=2. For the base case the complete equi-2-partite graph gives OE(n; n=2; 1. For the induction step, suppose we know that OE(n; d; a d-regular bipartite graph G(n; d; 2) exists, and by Proposition 5.2 it has a perfect matching M . Remove all edges in M from G to obtain a (d \Gamma 1)-regular bipartite graph G(n; d \Gamma Hence OE(n; d \Gamma Suppose n is odd. Then (5.1) implies that it is enough to show OE(n; d; Now d must be even by (2.3), and d - (n \Gamma 1)=2. Because by (5.5). Consider G := G(n \Gamma 5.2 we may find a perfect matching of G, say from G the edges and add to G a new vertex z and the edges fz; x i g and it is easy to see that the resulting graph is a d-regular 3-partite graph with n vertices, which proves (5.6). (b) Let G = G(n; d; r), which exists by hypothesis. Since d - n, Proposition 5.1 guarantees that G has a Hamilton cycle C. Removing all edges from C yields a G(n; d \Gamma 2; r), so OE(n; d \Gamma 1. If moreover n is even, then C has even length and we get a perfect matching M by taking alternate edges in C. Removing all edges in M from G yields a G(n; d \Gamma 1; r), so 1 in this case. \Xi Lemma 5.2. If r - 3, then provided that nd 1 and nd 2 are both even. Proof. Suppose are done. Suppose used inductively on decreasing d gives For odd n, since nd 1 and nd 2 are both even, both d 1 and d 2 must be even. Now Lemma 5.1(b) gives so (5.7) follows. \Xi Proof of Theorem 1.1. To commence the proof, we define p and q by that is r c - 1. Note that r divides only if In terms of p and q the assertions of the theorem then become: (i) If (iii) If 2. To prove (i)-(iii), we use the complete equi-r-partite graph G r (n) defined as follows. The graph G r (n) has vertices we define the vertex sets The edge set of G r (n) is E(G r Here is an equi-r-partition of V with For 1 - a - b - r we let G r a;b denote the induced subgraph of G r (n) on the vertex set To prove (i), if is an (n \Gamma p)-regular graph, hence Lemma 5.2 implies OE(n; d; To prove (ii), let q+1;r . Then (5.10) shows that H is a p(r \Gamma q \Gamma 1)-regular graph having vertices. Now r \Gamma q - 2 implies that H has degree p(r \Gamma q \Gamma 1), which is greater than half its vertices, so H has a Hamilton cycle C by Proposition 5.1. If p(r \Gamma q) is even, then H has a perfect matching M obtained by taking every other edge in C. Removing all edges in M from G r (n), the resulting graph is (n then completes the proof of (ii). If p(r \Gamma q) is odd, then p is odd, hence so is by (2.3). Thus it suffices to show that this case, for then Lemma 5.2 gives OE(n; d; 2. 1;q . Then H 0 is a (p vertices. If q ? 1 then hence H 0 is Hamiltonian. Since (p + 1)q is even, H 0 has a perfect matching M 0 . Removing all edges in M 0 [C from G r (n), the resulting graph is (n Suppose 1. Notice that since p(r \Gamma 1) is odd, r 6= 3, hence r - 4. Let H 00 be the induced subgraph of G r (n) on the set r and Then the number of vertices of H 00 is p(r and the minimum degree of H 00 is Proposition 5.1 implies that H 00 has a Hamiltonian cycle C 00 . By removing all edges in C 00 [E from G r (n) we have an (n To prove (iii) we proceed by induction on r, with an induction step from r to r 2. There are two base cases, Base Case r = 3. We have 2. Let Consider the graph G obtained by removing from G 3 (n) all edges in g. Then it is easy to see that G is (n gives OE(n; d; 2. Base Case 4. We have 3. Suppose first that p is odd. We relabel the vertices of G 4 (n) so that the sets X j in (5.9) become Let H be the subgraph of G 4 (n) induced on the vertex set fw 3g. Then even and H is Hamiltonian. Thus H has a perfect matching, call it M . Consider the edge set and form a graph G by removing all edges in E [ M from G 4 (n). Then G is an (n regular subgraph of G 4 (n), hence OE(n; Lemma 5.2. Suppose now that p is even. Then is forbidden by (2.3). It suffices therefore to show that OE(n; this case, for then Lemma 5.2 gives OE(n; d; 2. We use the vertex labelling (5.12), and let H be the subgraph of G 4 (n) induced on fw 3pg. Then jV so H has a perfect matching M . Consider the edge set Form a graph G by removing (n). It is an (n Induction Step. Fix r - 5 and define nd is eveng ; so 3. It is enough to show that OE(n; d yields OE(n; d; To do this, set furthermore we easily check that We may apply the induction hypothesis at r to conclude that there exists a d 0 -regular (r \Gamma 2)-partite graph a d 1 -regular bipartite graph with 2(p vertices disjoint from those of G; such a graph H exists by Lemma 5.1(a). Take the disjoint union of G and H and add in all edges between V (G) and V (H) to obtain a new graph G 0 on n vertices which is d 0 (n; r)-regular, according to (5.13). Thus OE(n; d completing the induction step for (iii). \Xi 6. Minimal Interference Level: Theorems 2.3 and 2.4 In this section we study the range fl - 3 and prove Theorems 2.3 and 2.4. The cases where '(n; ff; d) = 0, i.e. for d smaller than about follow from Theorem 1.1 applied with For the remaining cases, the harder step in the proofs is obtaining the (exact) lower bounds for '(n; ff; d). The upper bounds are obtained by explicit construction. We proceed to derive a lower bound for '(n; ff; d) stated as Lemma 6.2 below. Let G be any d-regular graph on n-vertices, let ng be a given coloring of G, and let ff also be given. We begin by partitioning the n colors into fl groups f ~ that each group ~ A i consists of consecutive colors and the groups ~ A fl are themselves consecutively arranged with respect to the cyclic ordering of colors (mod n), with all groups but ~ A 1 containing exactly ff A 1 contains the remaining ff m is given by and such a partition is completely determined by the choice of ~ We now choose ~ A 1 so as to minimize the number of vertices v in G that are assigned colors f(v) in ~ A 1 . After doing this, we have the freedom to cyclically relabel the colors (via the map affecting which edges have vertex colors that interfere. We use this freedom to specify that ~ in which case ~ see Figure 6.1. Notice that for 2 - i - fl any two colors in ~ A i interfere with each other. Figure 6.1 about here This partition of the colors induces a corresponding partition of the vertices of G into the color classes Now set a We now count the edges in G and in its complement - various ways. For any two subsets V and W of vertices, let e(V; W ) count the number of edges between vertices in V and those in W , and let - count the number of edges between A i and A j that are not in G, which is -a i;j := a i a Along with this we define a The d-regularity of G then yields a The potential interfering edge set B i;j between vertices in A i and those in A j is The actual interfering edge set is and we set c i;j := We clearly have - a i;j Finally, let ffi and ffi count the potential and actual non-interfering edges in A 1 , respectively, i.e. Certainly ffi - ffi. Since all edges between the vertices in the same component A i interfere, except for ffi edges in A 1 , we obtain the bound I ff G; To bound this further, we need the following bounds for edges connecting a vertex in the color set ~ A 1 to a vertex in its two neighboring color sets ~ A 2 and ~ A fl . Lemma 6.1. We have and Proof. We start with (6.6). By (6.4) it is enough to show that It suffices to show for fixed v 2 A 1 with ff because, using ff - m, this implies that, for sums over v 2 A 1 with ff To prove (6.8), given v 2 A 1 with ff we define the vertex set This is a set of ff by the minimizing property of the color set ~ A 1 . Now ff implies that Thus which is exactly (6.8). Thus (6.6) follows. The proof of (6.7) is analogous. \Xi To state the lower bound lemma, recall that the quantities p and q are defined by so Lemma 6.2. If d - Proof. We derive this result from the general bound I ff (G; f) - \Sigma fl where a for the vertex partition (6.2). To establish (6.10), we first note that Lemma 6.1 Together with (6.3), this yields Since the left side of this inequality is an integer, However, (6.3) also gives Substituting these bounds in (6.5) yields (6.10). To derive (6.9), we minimize the right side of (6.10) over all possible values: a i - 0 subject to \Sigma fl It is easy to verify that this occurs when all the a i 's are as equal as possible, i.e. 8 q of the a i take the value of the a i take the value Thus I ff (G; f) - qd(p which gives (6.9). \Xi Proof of Theorem 2.3. (a) This bound follows from Theorem 1.1, taking noting that (b) For d ? first establish the lower bounds where is odd or if n \Gamma d is even and p is even; even and p is odd ; using Lemma 6.2. The case simplifies to Now (6.12) follows on determining the cases for which p(d To show that this bound is attained, we simply construct the graph G with the coloring f that makes (6.11) hold. The constructions are easy and are left to the reader. \Xi Proof of Theorem 2.4. (a) The bounds where '(n; ff; d) = 0 follow from Theorem 1.1 with There remains the case in which (where Theorem 1.1 does not apply). We must show that For the upper bound ' - p, it suffices to construct an appropriate graph. Note first that p must be even since if p is odd then is also odd, contradicting (2.3). Now consider the equi-fl-partite graph G fl (n) defined in the proof of Theorem 1.1. We take a perfect matching M from the induced subgraph of G fl (n) on the vertex set (X fl We remove all the edges in M from G fl (n) and add the edges g. Then it is straightforward to check that the resulting graph G is (n 1)-regular and it clearly has exactly pinterfering edges when the sets X i are colored with fl mutually noninterfering colors. To show the lower bound ' - p G be an (n f an n-coloring of V (G) such that I ff (G; Take the partition fA associated to f constructed at the beginning of this section. We consider cases. Case (i). a 1 2. The minimality property of A 1 implies that, for all (G)nfvg such that jf(v) \Gamma f(w)j - ff. Thus I ff (G; f) - n=2 ? p=2. Case (ii). a Here the equality in (6.5) combined with I ff (G; f) - \Sigma fl Using (6.3) we then have I ff (G; f) - Case (iii). All a case requires that of the a i equal one a i equals p. Suppose first that a p. Observe that (6.14) and (6.3) yield I ff (G; f) - \Sigma fl Now (6.2) and a a Substituting this in (6.15) gives I ff (G; f) - pSuppose finally that a only one a a both. We treat only the case that a since the argument for a is similar. Let a i 0 p. Now by (6.5) and (6.3) I ff (G; f) - 1 a i Lemma 6.1 gives -a 1;2 I ff (G; f) - 1 completing case (iii). (b) We start from the formula (6.9) of Lemma 6.2, which gives a lower bound. We claim that equality occurs. This formula of '(n; ff; d) splits into several cases, according to when are integers or half-integers, and consideration of the parities of n \Gamma d and p leads to the formulas for ' in (2.12). For the upper bound, obtaining equality in the formula for '(n; ff; d) requires (6.11) to hold, and this easily determines the construction of a suitable graph G and a coloring f . We omit the details. \Xi 7. Minimax Interference Level: Theorems 2.5 and 2.6 We conclude by proving the bounds for L(n; ff; d) stated in Section 2. Proof of Theorem 2.5. To show part (a), the condition L(n; ff; d) certainly holds if the chromatic number -G - fl for all G 2 G(n; d). This holds for fl ? d by Brooks' theorem (Proposition 3.1). For the case we use the strong version of Brooks' theorem, which states that -G - \Delta G if no component of G is an odd cycle or a complete subgraph. Here d, and d - 3 implies there are no odd cycles, while the condition any connected component being the complete subgraph K d , for any other components must be d-regular but have at most d vertices, a contradiction. To show part (b), suppose that n - 2(d 1). Let G 2 G(n; d) consist of a complete graph K d+1 plus a d-regular graph G 0 on the other vertices. If d is odd then n is even, so that even, and the existence of G 0 is assured by a theorem of Erd-os and Gallai (1960) for simple graphs with specified degree sequences. If fl - d, at least two vertices of K d+1 interfere, so L ? 0. Suppose that 1. This implies d - 2a because we presume that 2. Let G consist of two disjoint copies of K d+1\Gammaa , adding edges between them that increase every degree to d. Each vertex requires a such edges, and this is feasible because a, at least two vertices of K d+1\Gammaa interfere, so L ? 0. Suppose finally that 1. Then n is odd, so d must be even. consist of two disjoint graphs G additions and deletions as follows. Add a edges from each G 1 vertex to G 2 vertices in as equal a way as possible for resulting vertex degrees in G 2 . Then each vertex in G 1 has degree d, x vertices in G 2 have degree d + 1, and y vertices in G 2 have degree d, where These equations imply that so x is even. We then remove x=2 edges within so that all vertices have degree d. We thus arrive at a graph G 2 G(n; d). If fl - d \Gamma a then at least two vertices in G 1 interfere, so L ? 0. Thus part (b) holds. \Xi Proof of Theorem 2.6. Suppose that 3 - d - W be nonnegative integers that satisfy To derive the upper bound on L in Theorem 2.6, let G be any graph in G(n; d). Let S denote the family of all partitions of the vertex set of G into fl groups, with q groups of size groups of size p. We adopt a probability model for S that assigns probability 1=jSj to each partition. Whichever partition obtains, we use fl mutually noninterfering colors for the fl groups in the partition. Suppose fu; vg is an edge in G. The probability that u and v lie in the same part of a member of S, so that fu; vg is an interference edge, is The expected number E[I ] of interference edges is nd=2 times this amount, i.e., so some member of S has a coloring that gives less than or equal to E[I ] edges whose vertices interfere. This is true for every G 2 G(n; d). Therefore we get the upper bound For the lower bound, assume initially that (d Let G consist of U disjoint copies of K d+1 . Then L(n; ff; d) - UL(d is the minimum number of interfering edges in K d+1 for an [n]. The analysis in Lemma 4.1 shows that L(d + 1; ff) is attained by an equi-fl-partition of V d+1 with f constant in each part. Since an equi-fl-partition of V d+1 has! Q groups of of P vertices each, we have To form G we begin with U disjoint copies of K d+1 and a disjoint KW . Each vertex in KW needs incident edges, so we add a total of W (d the K d+1 in such a way that W (d edges can be removed from within the K d+1 to end up with degree d for every vertex. Note that W (d and d would be odd. We ignore possible interference within KW and allow for the possibility that every edge removed from the K d+1 is an interference edge to get the lower bound --R On sublattices of the hexagonal lattice Graph Theory with Applications On coloring the nodes of a network Some theorems on abstract graphs Graphs with prescribed degrees of vertices (in Hungarian) The channel assignment problem for mutually adjacent sites Pair labellings of graphs theory and application On representatives of subsets Graph homomorphisms and the channel assignment problem Further results on T An extremal problem in graph theory (in Hungarian) --TR

graph coloring;interference threshold;regular graph

288859

Combinatorial Properties and Constructions of Traceability Schemes and Frameproof Codes.

In this paper, we investigate combinatorial properties and constructions of two recent topics of cryptographic interest, namely frameproof codes for digital fingerprinting and traceability schemes for broadcast encryption. We first give combinatorial descriptions of these two objects in terms of set systems and also discuss the Hamming distance of frameproof codes when viewed as error-correcting codes. From these descriptions, it is seen that existence of a c-traceability scheme implies the existence of a c-frameproof code. We then give several constructions of frameproof codes and traceability schemes by using combinatorial structures such as t-designs, packing designs, error-correcting codes, and perfect hash families. We also investigate embeddings of frameproof codes and traceability schemes, which allow a given scheme to be expanded at a later date to accommodate more users. Finally, we look briefly at bounds which establish necessary conditions for existence of these structures.

Introduction Traceability schemes for broadcast encryption were defined by Chor, Fiat and Naor [8], and frameproof codes for digital fingerprinting were proposed by Boneh and Shaw [4]. Although these two objects were designed for different purposes, they have some similar aspects. One of the purposes of this paper is to investigate the relations between traceability schemes and frameproof codes. We first give combinatorial descriptions of these two objects in terms of set systems, and also discuss the Hamming distance of frameproof codes when viewed as error-correcting codes. From these descriptions, it is seen that existence of a c-traceability scheme implies the existence of a c-frameproof code. In [4, 8], some constructions of frameproof codes and traceability schemes were pro- vided. We will provide new (explicit) constructions by using combinatorial structures such as t-designs, packing designs, error-correcting codes and perfect hash families. We also investigate embeddings of frameproof codes and traceability schemes, which allow a given scheme to be expanded at a later date to accommodate more users. Finally, we look briefly at bounds which establish necessary conditions for existence of these structures. In this rest of this section we review the definitions of c-frameproof codes and c- traceability schemes which were given in [4] and [8], respectively. 1.1 Frameproof codes In order to protect a product (such as computer software, for example), a distributor marks each copy with some codeword and then ships each user his data marked with that codeword (for some examples of how this might be done in practice, see [5]). This marking allows the distributor to detect any unauthorized copy and trace it back to the user. Since a marked object can be traced, the users will be deterred from releasing an unauthorized copy. However, a coalition of users may detect some of the marks, namely the ones where their copies differ. They can then change these marks arbitrarily. To prevent a group of users from "framing" another user, Boneh and Shaw [4] defined the concept of c-frameproof codes. A c-frameproof code has the property that no coalition of at most c users can frame a user not in the coalition. Let v and b be positive integers (b denotes the number of users in the scheme). A set is called a (v; b)-code and each w (i) is called a codeword. So a codeword is a binary v-tuple. We can use a b \Theta v matrix M to depict a (v; b)-code, in which each row of M is a codeword in \Gamma. \Gamma be a (v; b)-code. Suppose \Gamma. For we say that bit position i is undetectable for C if Let U(C) be the set of undetectable bit positions for C. Then is called the feasible set of C. (If then we define F .) The feasible set F (C) represents the set of all possible v-tuples that could be produced by the coalition C by comparing the d codewords they jointly hold. If there is a codeword w (j) 2 F (C)nC, then user j could be "framed" if the coalition C produces the v-tuple w (j) . The following definition from [4] is motivated by the desire for this situation not to occur. Definition 1.1 A (v; b)-code \Gamma is called a c-frameproof code if, for every W ' \Gamma such that We will say that \Gamma is a c-FPC(v; b) for short. Thus, in a c-frameproof code, the only codewords in the feasible set of a coalition of at most c users are the codewords of the members of the coalition. Hence, no coalition of at most c users can frame a user who is not in the coalition. Example 1.1 ([4]) For any integer b, there exists a b-FPC(b; b). The matrix depicting the code is a b \Theta b identity matrix. Example 1.2 There exists a 2-FPC(3; 4). The matrix depicting the code is as 1.2 Traceability schemes In many situations, such as a pay-per-view television broadcast, the data is only available to authorized users. To prevent an unauthorized user from accessing the data, the data supplier will encrypt the data and give the authorized users keys to decrypt it. Some unauthorized users (pirate users) might obtain some decryption keys from a group of one or more authorized users (called traitors). Then the pirate users can decrypt data that they are not entitled to. To prevent this, Chor, Fiat and Naor [8] devised a traitor tracing scheme, called a traceability scheme, which will reveal at least one traitor on the confiscation of a pirate decoder. Suppose there are a total of b users. The data supplier generates a base set T of v keys and assigns k keys to each user. These k keys comprise a user's personal key, and we will denote the personal key for user U by P (U ). A message consists of an enabling block and a cipher block. A cipher block is the encryption of the actual plaintext data using some secret S. The enabling block consists of data, which is encrypted using some or all of the v keys in the base set, the decryption of which will allow the recovery of S. Every authorized user should be able to recover S using his or her personal key, and then decrypt the cipher block using S to obtain the plaintext data. Some traitors may conspire and give an unauthorized user a "pirate decoder", F . The pirate decoder F will consist of k base keys, chosen from T , such that F ' [U2C P (U ), where C is the coalition of traitors. An unauthorized user may be able to decrypt S using a pirate decoder F . The goal of the data supplier is to assign keys to the users in such a way that when a pirate decoder is captured and the keys it possesses are examined, it should be possible to detect at least one traitor in the coalition C, provided that jCj - c (where c is a predetermined threshold). Traitor detection would be done by computing jF "P (U)j for all users U . If jF "P (U)j - users V 6= U , then U is defined to be an exposed user. Definition 1.2 Suppose any exposed user U is a member of the coalition C whenever a pirate decoder F is produced by C and jCj - c. Then the scheme is called a c-traceability scheme and it is denoted by c-TS(k; b; v). Let us now briefly discuss the difference between our scheme and that of [8]. In [8], nk for some integer n, and the set T of base keys is partitioned into k subsets S i , each of size n. We will denote S Each personal key P (U) is a transversal of (S contains exactly one key from each S i ). Suppose the secret key S is chosen from an abelian group G. To encrypt S, the data supplier splits G such that every share r i with each of the n keys in S i by computing t . The nk values t i;j comprise the enabling block. Each authorized user has one key from S i , so he or she can decrypt every r i , and thus compute S. In our definition, we do not require that each personal key be a transversal. A personal key can be made up of any selection of k base keys from the set T . The data supplier can use a k out of v threshold scheme (such as the Shamir scheme [13], for example) to construct v shares of the key S, and then encrypt each share r i with the key s i , for every s Note that our definition is a generalization of the one given in [8]. However, the generalization has to do with the way that the enabling block is formed, and not with the traceability property of the scheme. Our definition of the traceability property is the same as in [8]. Example 1.3 We present a 2-TS(5; 21; 21). The set of base keys is Z 21 . The personal key for user i (0 - i - 20) is where all arithmetic is done in Z 21 . (This is an application of a construction we will present in Theorem 3.5.) It can be shown that any two base keys occur together in exactly one personal key. Now, consider what happens when two traitors U and V construct a pirate decoder, F . The pirate decoder F must contain at least three personal keys from P (U) or However, for any other user W 6= U; V , 2. Hence either U or V will be the exposed user if the pirate decoder F is examined. 1.3 Previous results In the construction of frameproof codes and traceability schemes, the main goal is to accommodate as many users as possible. In other words, we want to find constructions with b as large as possible, given values for the parameters c and v (and k, in the case of traceability schemes). In general, we would prefer explicit constructions for these objects as opposed to non-constructive existence results. For example, Boneh and Shaw [4] proved the following interesting result. Theorem 1.1 For any integers c; v ? 0, there exists a c-FPC However, as noted in [4], the proof is not constructive. Hence, they also provide an explicit construction for a c-FPC Similarly, Chor, Fiat and Naor [8] gave an interesting non-constructive existence result for traceability schemes, as follows. Theorem 1.2 For any integers c; v ? 0, there exists a c-TS(v=(2c 2 We will provide several explicit constructions for frameproof codes and traceability schemes later in this paper. Although our constructions may not be as good asymptotically as those in [4] and [8], they will often be better for relatively small values of c and v. (For example, in order to obtain b - 2 in Theorem 1.1, it is necessary to take v - 16c 2 , so the construction is not useful for small values of v.) As well, our constructions are very simple and could be implemented very easily and efficiently. Combinatorial descriptions In this section, we give combinatorial descriptions of c-frameproof codes and c-traceability schemes. From these descriptions, it is fairly easy to see that the existence of a c-TS(k; b; v) implies the existence of a c-FPC(v; b). We will use the terminology of set systems. A set system is a pair (X; B) where X is a set of elements called points, and B is a set of subsets of X, the members of which are called blocks. A set system can be described by an incidence matrix. Let (X; B) be a set system g. The incidence matrix of (X; B) is the b \Theta v matrix ae Conversely, given an incidence matrix, we can define an associated set system in an obvious way. 2.1 Description of c-frameproof codes Since a c-FPC(v; b) is a b \Theta v (0; 1)-matrix, we can view a frameproof code as an incidence matrix or as a set system, as defined above. We have the following characterization of frameproof codes as set systems. Theorem 2.1 There exists a c-FPC(v; b) if and only if there exists a set system (X; B) such that and for any subset of d - c blocks there does not exist a block B 2 BnfB such that d d Proof. Suppose are d codewords in a c-FPC(v; b) (d - c). Without loss of generality, assume that in these codewords the first s bit positions are 0, the next t bit positions are 1, and in every other bit position at least one of the d codewords has the value 0 and at least one has the value 1. (Hence, the undetectable bit positions are the first s bit positions.) Then, it is not hard to see that the frameproof property is equivalent to saying that any other codeword w has at least one 1 in the first s bit positions, or at least one 0 in the next t bit positions. In other words, there does not exist a codeword with 0's in the first s bit positions and 1's in the next t bit positions. are the blocks in the set system corresponding to the d codewords d and d Hence the frameproof condition is equivalent to saying that there does not exist a block B such that "B 2.2 Description of c-traceability schemes Since a c-TS(k; b; v) consists of b k-subsets of a v-set, we can think of it as a set system, where X is the set of base keys and B is the set of personal keys. Theorem 2.2 There exists a c-TS(k; b; v) if and only if there exists a set system (X; B) such that B, with the property that for every choice of d - c blocks for any k-subset F ' [ d there does not exist a block B 2 BnfB such that Proof. Suppose (X; B) is a c-TS(k; b; v). For every set of d - c personal B, for any k-subset F ' [ d (i.e., a pirate decoder) and for any other personal key B, there exists a d) such that there is no block d. The converse is also straightforward. 2.3 Relationship of traceability schemes and frameproof codes We prove the following theorem relating traceability schemes and frameproof codes. Theorem 2.3 If there exists a c-TS(k; b; v), then there exists a c-FPC(v; b). Proof. Let (X; B) be the set system corresponding to a c-TS(k; b; v). We prove that (X; B) is a c-FPC(v; b). Suppose not; then there exist d - c blocks, B, and a block such that But this contradicts Theorem 2.2 (letting 2.4 Hamming distance of 2-frameproof codes Now we investigate some properties of the Hamming distance of c-frameproof codes. For any (v; b)-code, let d(x; y) denote the Hamming distance of two codewords x; y. Denote and Theorem 2.4 A (v; b)-code \Gamma is 2-frameproof if and only if for all i Proof. Let w (i) ; w (j) and w (h) be any three distinct codewords. Without loss of generality, assume that U(fw (i) ; w (j) so the first r bits of w (i) and w (j) are the same. We have that d(w (i) ; w (j) Hamming distance is a metric, we have that it will be the case that if and only if there is at least one bit position within the first r bit positions such that w (h) is different from w (i) and w (j) . But this is just the condition that the code is 2-frameproof (as stated in the proof of Theorem 2.1). The following result is an immediate corollary of the previous lemma. Corollary 2.5 A (v; b)-code \Gamma is 2-frameproof if d We give an example to illustrate the application of this corollary. In [6], a simple explicit construction is given for a (q; (q prime power q. Hence, for q ? 81, we see that d In fact, we have verified by computer that d for the codes produced by this construction for all odd prime powers q such that 31 - q - 79. Applying Corollary 2.5, we obtain the following result. Theorem 2.6 For any odd prime power q - 31, there exists a 2-FPC(q; (q 3 Constructions from combinatorial structures In this section, we will give some constructions of frameproof codes and traceability schemes from certain combinatorial designs, including t-designs, packing designs and orthogonal arrays. All the results on design theory that we require can be found in standard references such as the CRC Handbook of Combinatorial Designs [9]. 3.1 Constructions using t-designs First we give the definition of a t-design. Definition 3.1 A t-(v; k; -) design is a set system (X; B), where B, and every t-subset of X occurs in exactly - blocks in B. Note that, by simple counting, the number of the blocks in a t-(v; k; 1) design is . We will use t-(v; k; 1) designs to construct frameproof codes and traceability schemes, as described in the following theorems. Theorem 3.1 If there exists a t-(v; k; 1) design, then there exists a c-FPC(v; proof. Denote distinct blocks, and let g. If there exists a d, such that t. Since we have a t-design with Hence, for any B 2 BnfB g, we have that B 6' [ d . The t-design is a set system satisfying the conditions of Theorem 2.1, so the conclusion follows. Similarly, we can construct traceability schemes from t-(v; k; 1) designs; the value of c obtained is smaller, however. Theorem 3.2 If there exists a t-(v; k; 1) design, then there exists a c-TS(k; proof. Suppose there exists a t-(v; k; 1) design (X; B). Let d be d - c distinct blocks. g. If F ' [ d there exists a B i , d, such that c r On the other hand, since we have Hence, it follows that This shows that the t-design is a set system satisfying the conditions of Theorem 2.2, and the conclusion follows. There are many known results on existence and construction of t-(v; k; 1) designs for 3. On the other hand, no is known to exist for t - 6. However, known infinite classes of 2- and 3-designs provide some nice infinite classes of frameproof codes and traceability schemes. We illustrate with a few samples of typical results that can be obtained. First, for 3 - k - 5, a 2-(v; k; 1) design exists if and only if v j 1 or k mod [9, Chapter I.2]. Hence, we obtain the following. Theorem 3.3 There exist frameproof codes as follows: 1. There exists a 2-FPC(v; v(v \Gamma 1)=6) for all v j 2. There exists a 3-FPC(v; 3. There exists a 4-FPC(v; Similarly, we have the following theorem about the existence of 2-traceability schemes (note that to get c - 2 when Theorem 3.2, we need k - 5). Theorem 3.4 There exists a 2-TS(5; v(v \Gamma 1)=20; v), for all v j A design is known as a projective plane of order q; such a design exists whenever q is a prime power (see [9, Chapter VI.7]). In a projective plane we have so the frameproof codes obtained from it are not interesting (in view of Example 1.1, which does better). However, the traceability schemes will be of interest. Theorem 3.5 There exists a b p qc-TS(q 1), for all prime powers q. Example 1.3 is in fact obtained from the case Theorem 3.5. We give another class of examples derived from 3-(q 2 +1; q +1; 1) designs (these designs are called inversive planes and exist if q is a prime power; see [9, Chapter VI.7]). Theorem 3.6 For any prime power q, there exists a 3.2 Constructions using packing designs Another type of combinatorial design which can be used to construct frameproof codes and traceability schemes are packing designs. We give the definition as follows. Definition 3.2 A t-(v; k; -) packing design is a set system (X; B), where for every B 2 B, and every t-subset of X occurs in at most - blocks in B. Using the same argument as in the proof of Theorem 3.1, we have the following construction for frameproof codes. Theorem 3.7 If there exists a t-(v; k; 1) packing design having b blocks, then there exists a c-FPC(v; b), where Similarly, we have the following construction for traceability schemes, using the same argument as in the proof of Theorem 3.2. Theorem 3.8 If there exists a t-(v; k; 1) packing design having b blocks, then there exists a c-TS(k; b; v), where We mentioned previously that no t-(v; k; 1) designs are known to exist if v However, for any t, there are infinite classes of packing designs with a "large" number of blocks (i.e., close to ). These can be obtained from designs known as orthogonal arrays, which are defined as follows. Definition 3.3 An orthogonal array OA(t; k; s) is a k \Theta s t array, with entries from a set of s - 2 symbols, such that in any t rows, every t \Theta 1 column vector appears exactly once. It is easy to obtain a packing from an orthogonal array, as shown in the next lemma. Lemma 3.9 If there is an OA(t; k; s), then there is a t-(ks; k; 1) packing design that contains blocks. proof. Suppose that there is a OA(t; k; s) with entries from the set f0; 1g. For every column (y in the orthogonal array, define a block consist of the blocks thus constructed. It is easy to check that (X; B) is a t-(ks; k; 1) packing design. The following lemma ([9, Chapter VI.7]) provides infinite classes of orthogonal arrays, for any integer t. Lemma 3.10 If q is a prime power and t ! q, then there exists an OA(t; q hence a t- packing design with q t blocks exists. From Theorem 3.7 and Lemma 3.10, we obtain the following. Theorem 3.11 For any prime power q and any integer t ! q, there exists a In this construction, b - 2 2c (for frameproof codes) and b - 2 traceability schemes). Also, the resulting traceability schemes are of the "transversal type" considered in [8]. 3.3 Constructions using perfect hash families In this section, we present another method to construct frameproof codes, which uses a perfect hash family. Definition 3.4 An (n; m;w)-perfect hash family is a set of functions F such that f : ng ! f1; ng such that there exists at least one f 2 F such that f j X is one-to-one. When , an (n; m;w)-perfect hash family will be denoted by PHF(N ; n; m;w). Observe that a PHF(N ; n; m;w) can be depicted as an N \Theta n matrix with entries from having the property that in any w columns there exists at least one row such that the w entries in the given w columns are distinct. Results on perfect hash families can be found in numerous textbooks and papers. Mehlhorn [12] is a good textbook source; more recent constructions can be found in the papers [2] and [3]. The following theorem tells us how to use a perfect hash family to enlarge a frameproof code. Theorem 3.12 If there exists a PHF(N ; n; m; c + 1) and a c-FPC(v; m), then there exists a c-FPC(Nv; n). proof. be a c-FPC(v; m), and let F be a PHF(N ; n; m; c+1). be the (Nv; n)-code consisting of the n codewords means concatenation of strings. We will show that \Gamma 0 is a c- FPC(Nv;n). g. Recall that U(W ) is the set of undetectable bit positions of W . Assume that there exists a codeword u (i c+1 F is a PHF(N ; n; m; c + 1), there exists an h 2 F such that hj C is one-to-one, where g. Thus we have c different codewords w 1, such that w (h(i c+1 )) is in the feasible set of cg. This contradicts the fact that \Gamma is c-frameproof. In [4], the following construction of c-frameproof codes from error-correcting codes is given. Theorem 3.13 If there exists a c-FPC(v; q) and an (N; n) q-ary code with minimum Hamming distance d min ? there exists a c-FPC(vN; n). Alon [1] gave a construction of perfect hash families from error-correcting codes. We observe that if we use a perfect hash family constructed by Alon's method to obtain a c- frameproof code by applying Theorem 3.12, then the resulting code is essentially the same as the one constructed using Theorem 3.13. However, it is possible to use other constructions for perfect hash families to obtain new examples of frameproof codes. We provide one illustration now, which uses the following recursive construction from [3]. Lemma 3.14 Suppose there exists a PHF(N 1. Then there exists a PHF for any integer j - 1. Example 3.1 There exists a PHF(2; 5; 4; 3) as follows: Theorem 3.15 For any integer j - 1, there exists a 2-FPC(6 \Theta 4 proof. From Lemma 3.14 and Example 3.1, we obtain a PHF(2 \Theta 4 Combine this perfect hash family with the 2-FPC(3; 4) given in Example 1.2, and apply Theorem 3.12. 4 Embeddings In many cases the number of users of a scheme will increase after the system is set up. Initially, the data supplier will constuct a scheme that will accommodate a fixed number of users (which we denoted by b). If the number of users eventually surpasses b, we would like a simple method of extending the scheme which is "compatible" with the existing scheme. In the case of a traceability scheme, we do not want to change the personal keys already issued when the scheme is expanded. In the case of a frameproof code, we do not want to have to recall software that has already been sold. To solve this problem, we will introduce the concept of embedding frameproof codes and traceability schemes in larger ones. Definition 4.1 Let \Gamma be a c-FPC(v; b) and let \Gamma 0 be a c-FPC(v Suppose that, for every codeword w 2 \Gamma, there exists a codeword w such that the first bit positions of w 0 are the same as w, and the remaining v positions of w 0 are all 0's. Then we say that \Gamma is embedded into Initially, the distributor could use the code \Gamma to mark the products. When the number of users surpasses b, then codewords in \Gamma 0 n\Gamma are used. Note that the embedding property ensures that the codewords in \Gamma do not have to be changed when we proceed to the larger code. A similar definition can be given for traceability schemes. Definition 4.2 Let T be the set of v base keys of a c-TS(k; b; v), and let T 0 be the set of v 0 base keys of a c-TS(k; b Suppose that every personal key of the c-TS(k; b; v) is also a personal key of the c-TS(k; b we say that the first scheme is embedded into the second scheme. Note that the definition of embedding is even simpler if we consider the set system formulation of frameproof codes and traceability schemes. Namely, we say that (X; B) is embedded into Since t-designs and packing designs are set systems, the above definition of embedding applies. In fact, embeddings of combinatorial designs have been extensively studied, so we have a convenient method of constructing embeddible frameproof codes and traceability schemes. For example, in the case of 2-designs, we have the following result. Theorem 4.1 If there exists a 2-(v; k; 1) design that can be embedded into a 2-(v design, then there exists a that can be embedded into a \Xip that can be embedded into a \Xip We give a couple of illustrations of this idea. For necessary and sufficient conditions for embedding 2-(v; k; 1) designs into 2-(v designs are known, namely v j result is known as the "Doyen-Wilson Theorem" [9, Chapter I.4]; for III.1].) This provides a convenient way of embedding 2- and 3-frameproof codes into larger ones by application of Theorem 4.1. The following theorems are obtained. Theorem 4.2 For all v j 6 such that v 0 - 2v+1, there exists a 2-FPC(v; (v that can be embedded into a 2-FPC(v Theorem 4.3 For all v j exists a 3-FPC(v; (v that can be embedded into a 2-FPC(v Here is a small example to illustrate. Example 4.1 Given an embedding of a 2-(7; 3; 1) design into a 2-(15; 3; 1) design, a 2- FPC(7; 7) can be embedded into a 2-FPC(15; 35). The 35 codewords of the 2-FPC(15; 35) are given in Figure 1 (the first seven codewords, when restricted to the first seven bit positions, form the embedded 2-FPC(7; 7)). @ 1101000 00000000 1000000 00110000 0100000 00011000 1000000 00001010 0100000 10000100 0001000 10100000 1000000 01000100 0010000 10010000 0001000 01001000 1000000 10000001 A Figure 1: A 2-FPC(7; 7) embedded into a 2-FPC(15; 35). It is also well-known that for any prime power q and for any integers i - j, there exists a which can be embedded into a 2-(q j ; q; 1) design (in other words, the affine geometry AG(i; q) is a subgeometry of AG(j; q); see [9, Chapter VI.7]). The following result is obtained. Theorem 4.4 Let q be a prime power, and let i and j be positive integers such that i - j. Then there exists a which can be embedded into a can be embedded into a 5 Bounds In this section, we investigate necessary conditions for existence for frameproof codes and traceability schemes. These take the form of upper bounds on b, as a function of c and v (and k, in the case of traceability schemes). First we will give a bound for frameproof codes. be a c- FPC(v; b). Recall that U(C) denotes the set of undetectable bit positions for a subset denotes the feasible set of C. For 1 - d - c, let We begin by stating and proving a simple lemma. Lemma 5.1 Suppose is a c-FPC(v; b), and suppose t are as defined above. Then proof. Suppose t be such that g. Clearly U(C) ' U(C 0 ); however, since which contradicts Definition 1.1. The next result provides an upper bound on b which depends on t c\Gamma1 . Theorem 5.2 Suppose is a c-FPC(v; b), and suppose t are as defined above. Then l t c\Gamma1m proof. Let W ' \Gamma be chosen such that jW any codeword w (i) 2 \GammanW , let R It is easy to see that R i ' R for all which contradicts the fact that \Gamma is c-frameproof). In other words, the subsets R i constructed above form a Sperner family with respect to the ground set R. By Sperner's Theorem (see, for example [11, Theorem 6.3]), we see that dt 1, the result follows. The following bound on b is an immediate corollary. Corollary 5.3 If \Gamma is a c-FPC(v; b), then proof. From Lemma 5.1, we have that t 2. The conclusion follows from Theorem 5.2. Recall that Example 1.1 gave a construction for c-FPC(c; c), and we constructed a FPC(3; in Example 1.2. In both of these examples, the bound of Corollary 5.3 is met with equality. Now we turn our attention to traceability schemes, where we provide an upper bound on b. In [8], it was shown that b - v k=c if a c-TC(k; v; b) exists. We give a stronger bound, which is also based on the following observation made in [8]. Lemma 5.4 Suppose (X; B) is a c-TC(k; v; b). Then, for any subset of d - c blocks there does not exist a block B 2 BnfB such that B ' proof. The proof is essentially the same as the proof of Theorem 2.3. For obvious reasons, the collection of subsets B is called c-cover-free. Now, applying [10, Proposition 2.1], which gives an upper bound on the cardinality of a c-cover-free collection of sets, the following result is immediate. Theorem 5.5 If a c-TS(k; b; v) exists, then the following bound holds: c e. 6 Comments Further results on frameproof codes can be found in the PhD thesis of Yeow Meng Chee [7, Chapter 9]. Chee gives a probabilistic construction for 2-frameproof codes that improves upon results in [4], and provides efficient explicit constructions for frameproof codes using the idea of superimposed codes. Acknowledgements We thank the referee for several helpful comments. The authors' research is supported by NSF grant CCR-9402141. --R Explicit construction of exponential sized family of k-independent sets Some recursive constructions for perfect hash families Electronic marking and identification techniques to discourage document copying. On a class of constant weight codes CRC Handbook of Combinatorial Designs A Course in Combinatorics Data Structures and Algorithms How to share a secret --TR --CTR Sylvia Encheva , Grard Cohen, Frameproof codes against limited coalitions of pirates, Theoretical Computer Science, v.273 n.1-2, p.295-304, February 2002 Omer Berkman , Michal Parnas , Jii Sgall, Efficient dynamic traitor tracing, Proceedings of the eleventh annual ACM-SIAM symposium on Discrete algorithms, p.586-595, January 09-11, 2000, San Francisco, California, United States Yan Zhu , Wei Zou , Xinshan Zhu, Collusion secure convolutional fingerprinting information codes, Proceedings of the 2006 ACM Symposium on Information, computer and communications security, March 21-24, 2006, Taipei, Taiwan J. Lfvenberg, Binary Fingerprinting Codes, Designs, Codes and Cryptography, v.36 n.1, p.69-81, July 2005 Maura Paterson, Sliding-window dynamic frameproof codes, Designs, Codes and Cryptography, v.42 n.2, p.195-212, February 2007 Alexander Barg , Gregory Kabatiansky, A class of I.P.P. codes with efficient identification, Journal of Complexity, v.20 n.2-3, p.137-147, April/June 2004 Tran Van Trung , Sosina Martirosyan, On a Class of Traceability Codes, Designs, Codes and Cryptography, v.31 n.2, p.125-132, February 2004 G. Cohnen , S. Encheva , S. Litsyn , H. G. Schaathun, Intersecting codes and separating codes, Discrete Applied Mathematics, v.128 n.1, p.75-83, 15 May Wen-Guey Tzeng , Zhi-Jia Tzeng, A Public-Key Traitor Tracing Scheme with Revocation Using Dynamic Shares, Designs, Codes and Cryptography, v.35 n.1, p.47-61, April 2005 Satoshi Obana , Kaoru Kurosawa, Bounds and Combinatorial Structure of Yevgeniy Dodis , Nelly Fazio , Aggelos Kiayias , Moti Yung, Scalable public-key tracing and revoking, Proceedings of the twenty-second annual symposium on Principles of distributed computing, p.190-199, July 13-16, 2003, Boston, Massachusetts Tran Van Trung , Sosina Martirosyan, New Constructions for IPP Codes, Designs, Codes and Cryptography, v.35 n.2, p.227-239, May 2005 Yevgeniy Dodis , Nelly Fazio , Aggelos Kiayias , Moti Yung, Scalable public-key tracing and revoking, Distributed Computing, v.17 n.4, p.323-347, May 2005 X. Ma , R. Wei, On a Bound of Cover-Free Families, Designs, Codes and Cryptography, v.32 n.1-3, p.303-321, May-July 2004 Dan Boneh , Brent Waters, A fully collusion resistant broadcast, trace, and revoke system, Proceedings of the 13th ACM conference on Computer and communications security, October 30-November 03, 2006, Alexandria, Virginia, USA Dan Boneh , Brent Waters, A fully collusion resistant broadcast, trace, and revoke system, Proceedings of the 13th ACM conference on Computer and communications security, October 30-November 03, 2006, Alexandria, Virginia, USA Carlo Blundo , Paolo Darco , Alfredo De Santis , Massimiliano Listo, Design of Self-Healing Key Distribution Schemes, Designs, Codes and Cryptography, v.32 n.1-3, p.15-44, May-July 2004 Stelvio Cimato , Antonella Cresti , Paolo D'Arco, A unified model for unconditionally secure key distribution, Journal of Computer Security, v.14 n.1, p.45-64, January 2006 Charles J. Colbourn, Multiple access communications using combinatorial designs, Theoretical aspects of computer science: advanced lectures, Springer-Verlag New York, Inc., New York, NY, 2002

traceability scheme;hash family;frameproof code;t-design

288861

A Randomness-Rounds Tradeoff in Private Computation.

We study the role of randomness in multiparty private computations. In particular, we give several results that prove the existence of a randomness-rounds tradeoff in multiparty private computation of $\fxor$. We show that with a single random bit, $\Theta(n)$ rounds are necessary and sufficient to privately compute $\fxor$ of n input bits. With $d\ge 2$ random bits, $\Omega(\log n/ d)$ rounds are necessary, and $O(\log n/ \log d)$ are sufficient. More generally, we show that the private computation of a boolean function f, using $d\ge 2 $ random bits, requires $\Omega(\log S(f)/ d)$ rounds, where S(f) is the sensitivity of f. Using a single random bit, $\Omega(S(f))$ rounds are necessary.

Introduction A 1-private (or simply, private) protocol A for computing a function f is a protocol that allows possessing an individual secret input, x i , to compute the value of f(~x) in a way that no single player learns more about the initial inputs of other players than what is revealed by the value of f(~x) and its own input 1 . The players are assumed to be honest but curious. Namely, they all follow the prescribed protocol A but they could try to get additional information by considering the messages they receive during the execution of the protocol. Private computations in this setting were the subject of a considerable amount of An early version of this paper appeared in Advances in Cryptology, Proceedings of Crypto '94, Y. Desmedt, ed., Springer-Verlag, Lecture Notes in Computer Science, Vol. 839, pp. 397-410, 1994. y Dept. of Computer Science, Technion, Haifa, Israel. e-mail: [email protected]; http://www.cs.technion.ac.il/-eyalk; Work on this research was supported by the E. and J. Bishop Research Fund, and by the Fund for the Promotion of Research at the Technion. Part of this research was performed while the author was at Aiken Computation Laboratory, Harvard University, supported by research contracts ONR-N0001491-J-1981 and NSF-CCR-90-07677. z Dept. of Computer Science, Tel-Aviv University, Tel-Aviv, Israel. e-mail: [email protected] 1 In the literature a more general definition of t-privacy is given. The above definition is the case work, e.g., [BGW88, CCD88, BB89, CK89, K89, B89, FY92, CK92, CGK90, CGK92, KMO94]. One crucial ingredient in private protocols is the use of randomness. Quantifying the amount of randomness needed for computing functions privately is the subject of the present work. Randomness as a resource was extensively studied in the last decade. Methods for saving random bits range over pseudo-random generators [BM84, Y82, N90], techniques for re-cycling random bits [IZ89, CW89], sources of weak randomness [CG88, VV85, Z91], and constructions of different kinds of small probability spaces [NN90, AGHP90, S92, KM93, KM94, KK94] (which sometimes even allow to eliminate the use of randomness). A different direction of research is a quantitative study of the role of randomness in specific contexts, e.g., [RS89, KPU88, BGG90, CG90, BGS94, BSV94]. In this work, we initiate a quantitative study of randomness in private computations. We mainly concentrate on the specific task of computing the xor of n input bits. However, most of our results extend to any boolean function. The task of computing xor was the subject of previous research due to its being a basic linear operation and its relative simplicity [FY92, CK92]. It is known as a "folklore theorem" (and is not difficult to show) that private computation of xor cannot be carried out deterministically (for n - 3). On the other hand, with a single random bit such a computation becomes possible: At the first round player P n chooses a random bit r and sends to P 1 the bit x n \Phi r. Then, in round i xors its bit x i\Gamma1 with the message it received in the previous round, and sends the result to P i . Finally, P n xors the message it received with the random bit r. Both the correctness and privacy of this protocol are easy to verify. The main drawback of this protocol is that it takes n rounds. Another protocol for this task computes xor in two rounds but requires a linear number of random bits: In the first round each player P i chooses a random bit r i . Then, player P i sends x i \Phi r i to P 1 and r i to In the second round P 2 xors all the (random) bits it received in the first round and sends the result to P 1 which xors all the messages it received during the protocol to get the value of the function. Again, both the correctness and privacy of this protocol are not hard to verify. In this work we prove that there is a tradeoff between the amount of randomness and the number of rounds in private computations of xor. For example, we show that while with a single random bit \Theta(n) rounds are necessary and sufficient 2 , with two random bits O(log n) rounds suffice. 3 Namely, with a single additional random bit, the number of rounds is significantly reduced. Additional bits give a much more "modest" saving. More precisely, we prove that with bits O(log n= log d) rounds suffice and\Omega\Gammad/3 n=d) rounds are required. Our upper bound is achieved using a new method that enables us to use linear combinations of random bits again and again (while preserving the privacy). The lower bounds are proved using combinatorial arguments, and they are strong in the sense that they also apply to protocols that are allowed to precisely, dn=2e rounds. This upper bound is achieved by a slight modification of the first protocol above. Assume, for simplicity, that n is even. At the first round, player Pn sends xn \Phi r to player P 1 , and at the same time sends r to player Pn\Gamma1 . The players then continue as in the above protocol, forwarding messages in parallel until the two messages meet. More precisely, in round xors the message it received with its own input and sends it to player P i and player P n\Gamma(i\Gamma1) xors the message it received with its own input and sends it to player Pn\Gammai . In round n, player P nreceives two messages and can compute the value of the function by xoring the two messages together with its own input. 3 All logarithms are base 2, unless otherwise indicated. make errors, and that they actually show a lower bound on the expected number of rounds. We also show that if protocols are restricted to certain natural types (that include, in particular, the protocol that achieves the upper bound) we can even improve the lower bound and show that \Theta(log n= log d) rounds are necessary and sufficient. Our lower bound techniques apply not only to the xor function, but in fact give lower bounds on the number of rounds for any boolean function in terms of the sensitivity of the function. Namely, we prove that with d - 2 random bits \Omega\Gammats/ S(f)=d) rounds are necessary to privately compute a boolean function f , whose sensitivity is S(f ). With a single random bit rounds are necessary. The question whether private computations in general can be carried out in constant number of rounds was previously addressed [BB89, BFKR90]. In light of our results, a promising approach to investigate this question may be by proving that if a constant number of rounds is sufficient then a large number of random bits is required. Subsequent to our work, several other works were done regarding the amount of randomness in privacy. In particular, the amount of randomness required for computing the function xor t-privately, for t - 2, was studied in [BDPV95, KM96]; in [KOR96] it is shown that the boolean functions that can be computed privately with a constant number of random bits are exactly the functions that have linear size boolean circuits. The rest of the paper is organized as follows: In Section 2 we give some definitions. In Section 3 we give an upper bound on the number of rounds required to privately compute xor. In section 4 we give lower bounds on the number of rounds to privately compute a boolean function, in terms of its sensitivity. We conclude in Section 5 with lower bounds on the expected number of rounds in terms of the average sensitivity of the function being computed. The appendix contains the improved lower bounds for restricted types of protocols. Preliminaries We give here a description of the protocols we consider, and define the privacy property of protocols. More rigorous definitions of the protocols are given in Section 4.1. 1g be any boolean function. A set of n players P i (1 - i - n), each possessing a single input bit x i (known only to it), collaborate in a protocol to compute the value of f(~x). The protocol operates in rounds. In each round each player may toss some coins, and then sends messages to the other players (messages are sent over private channels so that other than the intended receiver no other player can listen to them). It then receives the messages sent to it by the other players. In addition, each player at a certain round chooses to output the value of the function. We assume that each player knows its serial number and the total number of players n. We may also assume that each player P i is provided with a read-only random tape R i from which it reads random bits (rather than toss coins). Each player P i receives during the execution of the protocol a sequence of messages. Let c i be a random variable of the communication string seen by player P i , and let C i be a particular communication string seen by P i . Informally, privacy with respect to player P i means that player anything (in particular, the inputs of the other players) from C i , except what is implied by its input bit, and the value of the function computed. Formally, Definition 1: (Privacy) A protocol A for computing a function f is private with respect to player P i if for any two input vectors ~x and ~y, such that sequence of messages C i , and for any random tape R i provided to P i , where the probability is over the random tapes of all other players. 3 Upper Bound This section presents a protocol which allows n players to use d - 2 random bits for computing xor privately. This protocol takes O(log d n) rounds. (For the case similar protocol that uses dn=2e rounds was already described in the introduction.) All arithmetic operations in this section are done modulo 2. Consider the following protocol (which we call the basic protocol): First organize the n players in a tree. The degree of the root of the tree is d + 1, and the degree of any other internal node is d (assume for simplicity that n is such that this forms a complete tree). The computation starts from the leaves and goes towards the root by sending messages (each of them of a single bit) as follows: Each leaf player P i sends its input bit x i to its parent in the tree. Each internal node, after receiving messages from its d children, sums them up (modulo 2) together with its input bit x i and sends the result to its parent. Finally, the root player sums up the d receives together with its input bit and the result is the output of the protocol. While a simple induction shows the correctness of this protocol, and it clearly runs in O(log d n) rounds, it is obvious that it does not maintain the required privacy. The second idea will be to "mask" each of the messages sent in the basic protocol by an appropriate random bit (constructed using the d random bits available), in a way that these masks will disappear at the end, and we will be left with the (un-masked) output. To do so we assign the nodes of the above tree vectors in GF [2 d ] as follows (the meaning of those vectors will become clear soon): Assign to the root the vector 0). The children of the root will be assigned d vectors such that the vectors in any d-size subset of them are linearly independent and the sum of all the vectors is (for example, the d unit vectors together with the satisfy these requirements). Finally, in a recursive way, given an internal node which is assigned a vector v, we assign to its d children d linearly independent vectors whose sum is v (note that in particular none of these vectors is the ~ 0 vector) 4 . We now show how to use the vectors we assigned to the nodes, so as to get a private protocol. We will assume that the random bits b are chosen by some external processor. We will 4 For example, such a collection of d vectors can be constructed as follows: Since v 6= ~ 0 there exists an index i such that v 1. The first d \Gamma 1 vectors will be the . The last vector will be Obviously the sum of these d vectors is v and they are linearly independent. later see that this assumption can be eliminated easily. Let v be the vector assigned to some player which is a leaf in the tree. We will give this player a single bit r the vector consisting of the d random bits, and the product is an inner product (modulo 2). The players will use the basic protocol, described above, with the modification that a player in a leaf also xors its message with the bit r v it received (the other players behave exactly as before). We claim that for every player P i , if in the basic protocol it sends the message m when the input vector is ~x, then in the modified protocol it sends the message m+ (v i \Delta b), where v i is the vector assigned to this player. The proof goes by induction: It is trivially true for the leaf players. For internal nodes the message is calculated by adding the input of the players to the sum of the incoming messages. Using, the induction hypothesis this sum is the message received from the k'th child in the basic protocol, and v k is the vector assigned to the k'th child. Since the construction is such that v i , the vector assigned to P i , satisfies v then a simple algebraic manipulation proves the induction step. In particular, since the root is assigned the vector its output equals the output of the basic protocol. Hence, the correctness follows. We now prove the privacy property of the protocol. The leaf players do not receive any message, hence there is nothing to prove. Let P j be an internal node in the tree. Denote by d the messages it receives. We claim that for every vector and for any input vector, we have where the probability is over the random choice of b (note that in this protocol the players do not make internal random choices). In other words, fix any specific input vector ~x, then for every vector w, there exists exactly one choice of values for b , such that the messages that P j receives, when the protocol is executed with input ~x, are the vector w. Denote by d the vectors corresponding to the d children of P j in the tree, and let the messages they have to send in the basic protocol given the input vector ~x. As claimed, for every d, the message that the k-th child sends in the modified protocol can be expressed as ~r). With this notation, for having s the following linear system has to be satisfied: 2 are linearly independent, this system has exactly one solution, as needed. As for the root player the same argument can be applied to any fixed d-size subset of the receives. This gives us that given any input vector ~x, for all d-size messages vectors ~ Now take two input vectors ~x and ~y such that x root = y root and such that by the correctness of the protocol, given a specific d-size messages-vector, the d 1'st message is the same for ~x and ~y. Thus the privacy property holds with respect to the root too. Finally, note that we assumed that the random choices were made by some external processor. However, we can let one of the leaf players randomly choose the bits b supply each of the leaf players with the appropriate bit r v . As the leaf players only send messages in the protocol, the special processor that selects the random bits gets no advantage. Note that if a player is non-honest it can easily prevent the other players from computing the correct output. However, it cannot get any additional information in the above protocol, since the only message each player gets after sending its own message is the value of the function. We have thus proved the following theorem: Theorem 1: The function xor on n input bits can be computed privately using d - 2 random bits in O(log n= log d) rounds. 4 Lower Bounds In this section we prove several lower bounds on the number of rounds required to privately compute a boolean function, given that the total number of random bits the players can toss is d. The lower bound is given in terms of the sensitivity of the function. In Section 4.1 we give some formal definitions. In Section 4.2 we introduce the notion of sensitivity and present a lemma, central to our proofs, about sensitivity of functions. The proof of the lower bound appears in Section 4.3. 4.1 Preliminaries We first give a formal definition for the protocols. A protocol operates in rounds. In each round each player P i , based on the value of its input bit x i , the values of the messages it received in previous rounds, and the values of the coins it tossed in previous rounds, tosses a certain number of additional coins, and sends messages to the other players. The values of these messages may depend on all of the above, including the coins just tossed by P i . The player may also choose to output the value of the function as calculated by itself (this is done only once by each player). Then, each player P i receives the messages sent to it by the other players. To define the protocol more formally, we give the following definition: Definition 2: (View) ffl A time-t partial view of player P i consists of its input bit x i , the messages it has received in the first t \Gamma 1 rounds, and the coins it tossed in the first t \Gamma 1 rounds. We denote it by ffl A time-t view of player P i consists of its input bit x i , the messages it has received in the first rounds, and the coins it tossed in the first t rounds. We denote it by V iew t Intuitively, the partial view of a player P i in round t determines how many coins (if at all) toss in round t. Then, its view (which includes those newly tossed coins) determines the messages P i will send in round t, and whether and which value it will output as the value of f . The formal definition of a protocol is given below: Definition 3: A protocol consists of a set of functions R k which determine how many coins are tossed by P i in round k, and a set of functions M k (where M is a finite domain of possible message values), which determine the message sent by P i to P j at round k. To quantify the amount of randomness used by a protocol we give the following definition: Definition 4: A d-random protocol is a protocol such that for any input assignment, the total number of coins tossed by all players in any execution is at most d. Note that the definition allows that in different executions different players will toss the coins. This may depend on both the input of the players, and previous coin tosses. Next, we define the correctness of a protocol. We usually consider protocols that are always correct; protocols that are allowed to err will be considered in Section 5.1. Definition 5: A protocol to compute a function f is a protocol such that for any input vector ~x and every i, player P i always correctly outputs the value of f(~x). It is sometime convenient to assume that each player P i is provided with a random tape R i , from which it reads random bits (rather than to assume that the player tosses random coins). The number of random coins tossed by player P i is thus the rightmost position of this tape that it reads. We thus denote by R i a specific random tape provided to player P i , and by ~ the vector or the random tapes of all the players denote the random variable for these tapes and vector of tapes). Note that if we fix ~ R, we obtain a deterministic protocol. Furthermore, V iew t i , for any i and t, is a function of the input assignment ~x, and the random tapes of the players. We can thus write it as V iew t denote by T i (~x; ~ R)) the round number in which player P i outputs its result given input assignment ~x and random tapes for the players ~ R. Definition (Rounds Complexity) An r-round protocol to compute a function f is a protocol to compute f such that for all i, ~x, ~ R, we have T i (~x; ~ For the purpose of our proofs, we slightly modify our view of the protocol in the following way. Fix an arbitrary binary encoding for the messages in M . We will consider a protocol where each player sends instead of a single message from M , a set of boolean messages that represent the binary encoding of the message to be sent in the original protocol. These messages are sent "in parallel" in the same round. Henceforth when we refer to messages we refer to these binary messages. Clearly, the number of rounds remains the same. 4.2 Sensitivity of Functions In this section we include some definitions related to functions f finite domain. Then, we present some useful properties related to these definitions. Definition 7: (Sensitivity) ffl For a binary vector Y , denote by Y (i) the binary vector obtained from Y by flipping the i'th entry. ffl A function f is sensitive to its i-th variable on assignment Y , if f(Y ) 6= f(Y (i) ). is the set of variables to which the function f is sensitive on assignment Y . ffl The sensitivity of a function f , denoted S(f), is S(f) 4 ffl The average sensitivity of a function f , denoted AS(f), is the average of jS f (Y )j. That is, Y 2f0;1g n jS f (Y )j. ffl The set of variables on which f depends, denoted D(f), is D(f) 4 )g. we say that f depends on its i-th variable. The following claim gives a lower bound on the degree of error if we evaluate a function f by means of another function g, in terms of their average sensitivities. We use this property in our proofs. 1: Consider any two functions f; for at most Proof: Consider the n-dimensional hypercube. An f -good edge is an edge ~y) such that f(~y). By the definitions, the number of f-good edges is exactly 2 n AS(f). Therefore, there are at least 2 n AS(f)\GammaAS(g)edges which are f-good but not g-good. For each such edge either f(~x) 6= g(~x) or f(~y) 6= g(~y). Since the degree of each vertex in the hypercube is n there must be at least 2 n \Delta AS(f)\GammaAS(g) inputs on which f and g do not agree. Next, we prove a lemma that bounds the growth of the sensitivity of a combination of func- tions. This lemma plays a central role in the proofs of our lower bounds, and any improvement on it will immediately improve our lower bounds. Lemma 2: Let be a set of m functions Assume j. Define the function F (Y ) 4 F assumes at most 2 d different values (different vectors), then the sensitivity of F is at most C \Delta 5 An obvious bound is S(F m. However, for reasons that will become clear soon we are interested in bounds which are independent of m. Proof: Let Y be the assignment on which F has the largest sensitivity, i.e. jS f (Y )j - jS f (Y 0 )j for any assignment Y 0 . Without loss of generality assume that F (Y Consider the set of neighbors of Y on which F has a value different than (the cardinality of this set is the sensitivity of F ). There are at most 2 d \Gamma 1 values of F attained on the assignments in this set. Consider one such value q 2 f0; 1g m . There is at least one index j such that q and since the sensitivity of f j is at most C, there can be at most C assignments Y (i) with the value q. We get that the total number of assignments Y (i) for which F has a value other than is at most C \Delta 4.3 Lower Bound on the Number of Rounds In this subsection we prove the following theorem. Theorem 3: Let A be an r-round d-random (d - 2) private protocol to compute a boolean function f . Then, r The first step of our proof uses the d-randomness property of the protocol to show that the number of views a player can see on a fixed input ~x is at most 2 d (over the different random tapes of all the players). Note that this is not obvious; although only d coins are tossed during every execution, the identity of the players that toss these coins may depend on the outcome of previous coin tosses. Lemma 4: Consider a private d-random protocol to compute a boolean function f . Fix an input ~x. Let C k be the communication string seen by player P i up to round k on input ~x and vector of random tapes ~r. Then, for every player P i , C k can assume at most 2 d different values (over the different vectors of random tapes ~r). Proof: For each execution we can order the coin tosses (i.e., readings from the local random tapes) according to the rounds of the protocol and within each round according to the index of the players that toss them. The identity of the player to toss the first coin is fixed by ~x. The identity of the player to toss any next coin is determined by ~x, and the outcome of the previous coins. Therefore, the different executions on input ~x can be described using the following binary tree: In each node of the tree we have a name of a player P j that tosses a coin. The two outgoing edges from this node, labeled 0 and 1 according to the outcome of the coin, lead to two nodes labeled P k and P ' respectively (k; ', and j need not be distinct) which is the identity of the player to toss the next coin. If no additional coin toss occurs, the node is labeled "nil"; there are no outgoing edges from a nil node. By the d-randomness property of the protocol, the depth of the above tree is at most d, hence it has at most 2 d root-to-leaf paths. Every possible run of the protocol is described by one root-to-leaf path. Such a path determines all the messages sent in the protocol, which player tosses coins and when, and the outcome of these coins. In particular each path determines for any P i the value of C k k). Hence, C k at most 2 d different values. In the following proof we restrict our attention to a specific deterministic protocols derived from the original protocol by fixing specific vector of random tapes ~ players. In such a deterministic protocol the views of the players are functions of only the input assignment ~x. Lemma 5: Consider a private d-random protocol to compute a boolean function f . Fix random tapes ~ is the view of player P i at round k on input ~x and vector of random tapes ~r. Then, for any P i , V iew k R) can assume at most 2 d+2 different values (over the values of ~y). Proof: Partition the input assignments ~x into 4 groups according to the value of x i (0 or 1), and the value of f(~x) (0 or 1). We argue that the number of different values the view can assume within each such group is at most 2 d . Fix an input ~x in one of these 4 groups and consider any other input ~y pertaining to the same group. Recall that C k R) is the communication string seen by player P i until round k on input ~y and when the random tapes of the players are ~ R. If the value of C k R) is some communication string C i , then by the privacy requirement 6 , communication C i must also occur by round k when the input is ~x, and the vector of random tapes is some ~ (R Thus, the value of C k R) must also appear as some vector of random tapes. However, by Lemma 4, C k can assume at most values (over the values of ~r). Thus, C k R) can assume at most 2 d values over the possible input assignments that pertain to the same group. Now, observe that V iew k determined by the input bit y i , the communication string and the random tape r i . Therefore, on ~ R and on two input assignments ~y and ~y 0 of the same group (in particular y R) then C k R). Thus, V iew k R) can assume at most 2 d different values over the input assignments that pertain to the same group. The following lemma gives an upper bound on the sensitivity of the view of a player at a given round, in terms of the number of random bits and the round number. This will enable us to give a lower bound on the number of necessary rounds. Lemma Consider a private d-random protocol to compute a boolean function f , and consider specific vector of random tapes ~ R, and the deterministic protocol derived by it. Then for every player R) (as a function of ~x only) has sensitivity of at most Q(k) 4 Proof: First note that since we fix the random tapes, the views of the players are functions of the input assignment ~x only. We prove the lemma by induction. For the view of any player depends only on its single input bit. Thus, the claim is obvious. For k ? 1 assume the claim holds for any ' ! k. This implies in particular that all messages received by player P i and included in the view under consideration have sensitivity of at most Q(k \Gamma 1). Clearly the 6 The privacy requirement is defined on the final communication string, but this clearly implies the same requirement on any prefix of it. input bit itself has sensitivity 1 which is at most Q(k \Gamma 1). Thus the view under consideration is composed of bits each having sensitivity at most Q(k \Gamma 1). Moreover, by Lemma 5 the view can assume at most 2 d+2 values. It follows from Lemma 2 that the sensitivity of the view under consideration is at most Q(k Q(k). (Note that Lemma 2 allows us to give a bound which does not depend on the number of messages received by P i .) We can now give the lower bound on the number of rounds, in terms of the sensitivity of the function and the number of random bits. Theorem 7: Given a private d-random protocol (d - 2) to compute a boolean function f , consider the deterministic protocol derived from it by any given random tapes ~ R. For any player there is at least one input assignment ~x such that T i (~x; ~ Proof: Consider a fixed but arbitrary player P i . Denote by t the largest round number in which outputs a value, i.e., R)g. We claim that as long as the sensitivity of the view of P i does not reach S(f ), there is at least one input assignment for which P i cannot output the correct value of f . Let Y be an input assignment on which the sensitivity S(f) is obtained. That is, the value of F (Y ) is different than the value of F on S(f) of Y 's ``neighbors''. Hence, if the sensitivity of the view of P i is less than S(f ), then the output of P i must be wrong on either Y or on at least one of these "neighbors" (as the sensitivity of the view is an upper bound on the sensitivity of the output). Thus, t is such that S(V iew t 6, we get 2 (d+2)(t\Gamma1) - S(f ), i.e., t - log S(f) 1. This proves Theorem 3; moreover, it shows not only that there is an input assignment ~x and random tapes ~ R for which the protocol runs "for a long time", but also that for each vector of random tapes ~ R there is such input assignment. The following corollary follows for the function xor (using the fact that Corollary 8: Let A be an r-round d-random private protocol (d - 2) to compute xor of n bits. 4.3.1 Lower Bound for a Single Random Bit For the case of a single random bit (d = 1), we have the following lower bound: Theorem 9: Let A be an r-round 1-random private protocol to compute a boolean function f . To prove the theorem, we restrict our attention to one of the two deterministic protocols derived from the original protocol by a fixing the value of the random bit 7 . The messages and views in this protocol are functions of the input vector, ~x, only. Let Y be an assignment on which S(f ), the sensitivity of f , is obtained. For a given function m, a variables x j is called good for m on Y if both m and f are sensitive to x j on Y . We denote by Gm (Y ) the set of good variables on Y , i.e., Gm (Y ) 4 first prove the following two lemmas: 7 We let the identity of the player that tosses this coin to possibly depend on the input ~x. However note that if we want that the privacy and 1-randomness properties hold, this cannot be the case. Lemma 10: Consider any player P i . Denote by m 1 a message that P i receives such that 1. Then for any other message m 2 received by P i such that jG m 2 either (a) Gm 1 or (b) jG m 1 Proof: Assume towards a contradiction, that both (a) and (b) do not hold. First, since are two variables x k 2 Gm 1 such that k 6= i and ' 6= i. Moreover, by the assumption that (a) does not hold, we can assume without loss of generality (as to the names of m 1 and By the assumption that (b) does not hold, there is a variable x j (j 6= i) such that f is sensitive to x j on Y , but both m 1 and m 2 are not sensitive to x j on Y . Now consider the following three input assignments: Consider V iew i on the above 3 inputs and assume, without loss of generality, that m 1 (Y are not sensitive to x j on Y , then m 1 (Y (j) sensitive to x k on Y , then m 1 (Y sensitive to x ' on Y , but m 1 is not, then m 1 (Y (') different values for Y 0 , Y 1 , and Y 2 . The function f is sensitive on Y to all of j; k and ', therefore, is equal in all three assignments. However, in the proof of Lemma 5, it is shown that the number of values of V iew i corresponding to inputs with the same value of f and the same value of x i is at most 2 The following lemma gives an upper bound on the sensitivity of the view of the player in terms of the round number. We then use this lemma to give a lower bound on the number of necessary rounds. Lemma 11: Let t - (S(f)\Gamma1)=2 be a round number and P i be any player. Then, jG V iew t t. Proof: We prove the claim by induction on t. For getting any messages, the view depends only on x i ). For assume the claim holds for any k ! t. Denote by M the set of messages received by P i and included in the view under consideration. Clearly G V iew t There could be one of three cases: 1. For any message In this case the claim clearly holds. 2. Any two messages , such that jG m 1 g. It follows that jG V iew t by the induction hypothesis jG t, then jG V iew t t. 3. There are two messages , such that jG m 1 but Gm 1 g. By Lemma 10, jG m 1 and (without loss of This contradicts the induction hypothesis as received in some round k ! t - (S(f) \Gamma 1)=2, and therefore generated by a view of round k. By the induction hypothesis jG We can now give the proof of Theorem 9. Proof of Theorem 9. Consider any player P i . Denote by t the largest round number in which outputs a value, i.e. 0)g. As in the proof of Theorem 7, it must be that For the function xor we have the following corollary. Corollary 12: Let A be an r-round 1-random private protocol to compute xor of n bits. Then 5 Lower Bounds on the Expected Number of Rounds As the protocols we consider are randomized, it is possible that for the same input ~x, different random tapes for the players will result in executions that run for different number of rounds. Hence, it is natural to consider not only the worst case running time but also the expected running time. Usually, saying that a protocol has expected running time r means that for every input ~x the expected time until all players finish the execution is bounded by r (where the expectation is over the choices of the random tapes of the players). Here we consider a weaker definition, which requires only the existence of a player i whose expected running time is bounded by r. As we are proving a lower bound, this only makes our result stronger: It would mean that for every player there is an input assignment for which the expected running time is high. Note that it is not necessarily the case that the first player that computes the value of the function can announce this value (and thus all players compute the value within one round). The reason is that the fact that a certain player computes the function at a certain round may reveal some information on the inputs, and hence such announcing may violate the privacy requirement (see [CGK90]). We first define the expected rounds complexity of a protocol. Definition 8: (Expected Rounds Complexity) An expected r-round protocol to compute a function f is a protocol to compute f such that there exists a player P i such that for all ~x, The lower bound that we prove in this section is in terms of the average sensitivity of the computed function. In particular, we prove an\Omega\Gamma/28 n=d) lower bound on the expected number of rounds required by protocols that privately compute xor of n bits. We will prove the following theorem: Theorem 13: Let f be a boolean function and let A be an expected r-round d-random private protocol (d - 2) to compute the function f . Then, To prove the theorem we consider a protocol A and fix any player P i . We say that the protocol is late on input ~x and vector of random tapes ~ 1. We define to be 1 if and only if the protocol is late on ~x and ~r. For the purpose of our proofs in this section we define a uniform distribution on the 2 n input assignments (this is not to say that the input are actually drawn by such distribution). Moreover, note that the domain of vectors of random tapes is enumerable. We first show that for any deterministic protocol derived from a private protocol to compute f , not only there is at least one input on which the protocol is late, but that this happens for a large fraction of the inputs. Lemma 14: Consider a player P i and any fixed vector of random tapes ~ Then AS(f) Proof: Consider the views of P given the vector of random tapes ~ R. For any round t such that t ! log AS(f) function g computed from such a view can have at most the same sensitivity, and thus clearly an average sensitivity of at most AS(f). By Claim 1, such a function g can have the correct value for the function f for at most 2 n AS(f) assignments. Since we assume that A is correct for all input assignments, it follows that at least 2 n AS(f)\Gamma AS(f) input assignments are late. We can now give a lower bound on the expected number of rounds. Lemma 15: Consider a player P i . There is at least one input assignment ~x for which AS(f) log AS(f) Proof: By Lemma 14, E ~r;~x [L(~x; ~r)] - AS(f)\Gamma AS(f) . Hence, there is at least one input assignment ~x for which E ~r [L(~x; ~r)] - AS(f)\Gamma AS(f) . For such ~x we get AS(f) log AS(f) as needed. Theorem 13 follows from the above lemma. The following corollary applies to the function xor. Corollary 16: Let A be an expected r-round d-random private protocol (d - 2) to compute xor of n bits. Then, r Proof: Follows from Theorem 13 and the fact that 5.1 Weakly Correct Protocols In this section we consider protocols that are allowed to make a certain amount of errors. Given a protocol A, denote by A i (~x; ~r) the output of the protocol in player P i , given input assignment ~x and vector of random tape Definition 9: For ffi)-correct protocol to compute a function f is a protocol that for every player P i and every input vector ~x satisfies P r ~r [A i (~x; Note that while designing a protocol one usually wants a stronger requirement; that is, with high probability all players compute the correct value. With the above definition, it is possible that in every execution of the protocol at least one of the players is wrong. However, as our aim now is to prove a lower bound this weak definition only makes our result stronger. In the following theorem we give lower bounds on the number of rounds and on the expected number of rounds for weakly correct protocols. Theorem 17: Let f be a boolean function. ffl Let A be a (1 \Gamma ffi)-correct r-round d-random private protocol (d - 2) to compute f . If AS(f) then r AS(f)=d). ffl Let A be a expected r-round d-random private protocol (d - 2) to compute f . Then AS(f) Proof: We first prove the lower bound on the number of rounds, and then turn our attention to the expected number of rounds. The correctness requirement implies that for any player P i , This implies that there exists a vector of random tapes ~ R such that for at least 2 n (1 \Gamma ffi) input assignments ~x, A i (~x; ~ As in the proof of Lemma 14 (using Claim 1), it follows that before round number log AS(f) 1, the protocol can be correct on at most 2 n AS(f) inputs (with random tapes ~ R). Since we require that at least are correct, we have that at least AS(f) AS(f) inputs are late. To get a lower bound on r for an r-round protocol, it is sufficient to have a single input vector ~x such that the execution on (~x; ~ R) is "long". For this, note that if AS(f) then (for random tapes ~ R) the number of late inputs is greater than 0. This gives us a lower bound of r = \Omega\Gamma112 AS(f)=d) for any (1 \Gamma ffi)-correct r-round d-random protocol, with ffi as above. We now turn to the lower bound on the expected number of rounds of (1 \Gamma ffi)-correct protocols. Consider a player P i . Define a to be 1 if and only if A i (~x; Then, the correctness requirement implies that E ~r [G(~x; ~r)] all ~x. It follows that for any ~ R the probability that ~ ffiis at least 1 \Gamma 2ffi. 8 For any such vector of random tapes, ~ R, consider the deterministic protocol derived from it. In such a protocol there are at least s AS(f) AS(f) s A late input assignments; that is, E ~x [L(~x; ~ AS(f) ffi). Thus AS(f) s It follows that there is at least one input assignment ~x for which AS(f) s which implies that AS(f) s A \Delta log AS(f) d as claimed. The following gives the lower bounds for the function xor. Corollary 18: For fixed A be a (1 \Gamma ffi)-correct d-random expected r-round private protocol to compute xor of n bits. Then r n=d). (Obviously the same lower bound holds for r-round protocols.) Proof: Follows from Theorem 17 and the fact that n. Note that the expression ) is greater than 0 for any ffi ! 1=2 (and sufficiently large n). . Thus there is at least one input assignment ~x such that E ~r [G(~x; which is a contradiction to the protocol being correct. 6 Conclusion In this paper we initiate the quantitative study of randomness in private computations. As mentioned in the introduction, our work was already followed by additional work on this topic [BDPV95, KM96, KOR96]. We give upper bounds and lower bounds on the number of rounds required for computing xor privately with a given number of random bits. Alternatively, we give bounds on the number of random bits required for computing xor privately within a given number of rounds. Our lower bounds extend to other functions in terms of their sensitivity (and average sensitivity). An obvious open problem is to close the gap between the upper bound and the lower bound for computing xor using d random bits. One possible way of doing this is to improve the bound given by Lemma 2. Acknowledgments We thank G'abor Tardos for improving the constant and simplifying the proof of Lemma 2, and Demetrios Achlioptas for his help in simplifying the proof of Lemma 19. We also thank Benny Chor for useful comments. --R "Simple constructions of almost k-wise independent random variables" "Non-Cryptographic Fault-Tolerant Computing in a Constant Number of Rounds" "Perfect Privacy for Two-Party Protocols" "Security with Low Communication Overhead" "Randomness in Interactive Proofs" "Completeness Theorems for Non-Cryptographic Fault-Tolerant Distributed Computation" "How to Generate Cryptographically Strong Sequences Of Pseudo-Random Bits" "On the Dealer's Randomness Required in Secret Sharing Schemes" "Randomness in Distribution Protocols" "On the Number of Random Bits in Totally Private Computations" "Bounds on Tradeoffs between Randomness and Communication Complexity" "Multiparty Unconditionally Secure Protocols" "A Zero-One Law for Boolean Privacy" "A Communication-Privacy Tradeoff for Modular Addition" "Unbiased Bits from Sources of Weak Randomness and Probabilistic Communication Complexity" "Dispersers, Deterministic Amplification, and Weak Random Sources" "Communication Complexity of Secure Computation" "How to Recycle Random Bits" "(De)randomized Construction of Small Sample Spaces in NC" "Constructing Small Sample Spaces Satisfying Given Con- straints" "On Construction of k-wise Independent Random Variables" "A Time-Randomness Tradeoff for Oblivious Routing" "Privacy and Communication Complexity" "Randomness in Private Computations" "Reducibility and Completeness in Multi-Party Private Computations" "Characterizing Linear Size Circuits in Terms of Privacy" "Small-Bias Probability Spaces: Efficient Constructions and Appli- cations" "Pseudorandom Generator for Space Bounded Computation" "Memory vs. Randomization in On-Line Algorithms" "Sample Spaces Uniform on Neighborhoods" "Random Polynomial Time is Equal to Slightly-Random Polynomial Time" "Theory and Applications of Trapdoor Functions" "Simulating BPP Using a General Weak Random Source" --TR --CTR Balogh , Jnos A. Csirik , Yuval Ishai , Eyal Kushilevitz, Private computation using a PEZ dispenser, Theoretical Computer Science, v.306 n.1-3, p.69-84, 5 September Anna Gal , Adi Rosen, Lower bounds on the amount of randomness in private computation, Proceedings of the thirty-fifth annual ACM symposium on Theory of computing, June 09-11, 2003, San Diego, CA, USA Anna Gl , Adi Rosn, A theorem on sensitivity and applications in private computation, Proceedings of the thirty-first annual ACM symposium on Theory of computing, p.348-357, May 01-04, 1999, Atlanta, Georgia, United States Eyal Kushilevitz , Rafail Ostrovsky , Adi Rosn, Amortizing randomness in private multiparty computations, Proceedings of the seventeenth annual ACM symposium on Principles of distributed computing, p.81-90, June 28-July 02, 1998, Puerto Vallarta, Mexico

lower bounds;randomness;private distributed computations;sensitivity

288866

The Number of Intersection Points Made by the Diagonals of a Regular Polygon.

We give a formula for the number of interior intersection points made by the diagonals of a regular n-gon. The answer is a polynomial on each residue class modulo 2520. We also compute the number of regions formed by the diagonals, by using Euler's formula 2.

Introduction We will find a formula for the number I(n) of intersection points formed inside a regular n-gon by its diagonals. The case depicted in Figure 1. For a generic convex n-gon, the answer would be , because every four vertices would be the endpoints of a unique pair of intersecting diagonals. But I(n) can be less, because in a regular n-gon it may happen that three or more diagonals meet at an interior point, and then some of the intersection points will coincide. In fact, if n is even and at least 6, I(n) will always be less than , because there will be diagonals meeting at the center point. It will result from our analysis that 4, the maximum number of diagonals of the regular n-gon that meet at a point other than the center is 3 if n is even but not divisible by 6; 5 if n is divisible by 6 but not 30, and; 7 if n is divisible by 30: with two exceptions: this number is 2 if In particular, it is impossible to have 8 or more diagonals of a regular n-gon meeting at a point other than the center. Also, by our earlier remarks, the fact that no three diagonals meet when n is odd will imply that for odd n. Date: January 30, 1995. 1991 Mathematics Subject Classification. Primary 51M04; Secondary 11R18. Key words and phrases. regular polygons, diagonals, intersection points, roots of unity, adventi- tious quadrangles. The first author is supported by an NSF Mathematical Sciences Postdoctoral Research Fellowship. Part of this work was done at MSRI, where research is supported in part by NSF grant DMS-9022140. Figure 1. The 30-gon with its diagonals. There are 16801 interior intersection points: 13800 two line intersections, 2250 three line inter- sections, 420 four line intersections, 180 five line intersections, 120 six line intersections, seven line intersections, and 1 fifteen line intersection DIAGONALS OF A REGULAR POLYGON 3 A careful analysis of the possible configurations of three diagonals meeting will provide enough information to permit us in theory to deduce a formula for I(n). But because the explicit description of these configurations is so complex, our strategy will be instead to use this information to deduce only the form of the answer, and then to compute the answer for enough small n that we can determine the result precisely. In order to write the answer in a reasonable form, we define Theorem 1. For n - 3, Further analysis, involving Euler's formula the number R(n) of regions that the diagonals cut the n-gon into. Theorem 2. For n - 3, These problems have been studied by many authors before, but this is apparently the first time the correct formulas have been obtained. The Dutch mathematician Gerrit Bol [1] gave a complete solution in 1936, except that a few of the coefficients in his formulas are wrong. misprints and omissions in Bol's paper are mentioned in [11].) The approaches used by us and Bol are similar in many ways. One difference (which is not too substantial) is that we work as much as possible with roots of unity whereas Bol tended to use more trigonometry (integer relations between sines of rational multiples of -). Also, we relegate much of the work to the computer, whereas Bol had to enumerate the many cases by hand. The task is so formidable that it is amazing to us that Bol was able to complete it, and at the same time not so surprising that it would contain a few errors! Bol's work was largely forgotten. In fact, even we were not aware of his paper until after deriving the formulas ourselves. Many other authors in the interim solved special cases of the problem. Steinhaus [14] posed the problem of showing that no three diagonals meet internally when n is prime, and this was solved by Croft and Fowler [3]. (Steinhaus also mentions this in [13], which includes a picture of the 23-gon and its diagonals.) In the 1960s, Heineken [6] gave a delightful argument which generalized this to all odd n, and later he [7] and Harborth [4] independently enumerated all three-diagonal intersections for n not divisible by 6. The classification of three-diagonal intersections also solves Colin Tripp's problem [15] of enumerating "adventitious quadrilaterals," those convex quadrilaterals for which the angles formed by sides and diagonals are all rational multiples of -. See Rigby's paper [11] or the summary [10] for details. Rigby, who was aware of Bol's work, mentions that Monsky and Pleasants also each independently classified all three-diagonal intersections of regular n-gons. Rigby's papers partially solve Tripp's further problem of proving the existence of all adventitious quadrangles using only elementary geometry; i.e., without resorting to trigonometry. All the questions so far have been in the Euclidean plane. What happens if we count the interior intersections made the diagonals of a hyperbolic regular n-gon? The answers are exactly the same, as pointed out in [11], because if we use Bel- trami's representation of points of the hyperbolic plane by points inside a circle in the Euclidean plane, we can assume that the center of the hyperbolic n-gon corresponds to the center of the circle, and then the hyperbolic n-gon with its diagonals looks in the model exactly like a Euclidean regular n-gon with its diagonals. It is equally easy to see that the answers will be the same in elliptic geometry. 2. When do three diagonals meet? We now begin our derivations of the formulas for I(n) and R(n). The first step will be to find a criterion for the concurrency of three diagonals. Let A; B; C; D; E; F be six distinct points in order on a unit circle dividing up the circumference into arc lengths u; x; v; z and assume that the three chords AD;BE;CF meet at P (see Figure 2). By similar triangles, Multiplying these together yields (AF and so Conversely, suppose six distinct points A; B; C; D; E; F partition the circumference of a unit circle into arc lengths u; x; v; z and suppose that (1) holds. Then the three diagonals AD;BE;CF meet in a single point which we see as follows. Let lines DIAGONALS OF A REGULAR POLYGON 5 z y A x F Figure 2. AD and BE intersect at P 0 . Form the line through F and P 0 and let C 0 be the other intersection point of FP 0 with the circle. This partitions the circumference into arc lengths As shown above, we have and since we are assuming that (1) holds for u; x; v; z we get sin(y 0 =2) Substituting above we get sin(y 0 =2) and so cot(y Now . Thus, the three diagonals AD;BE;CF meet at a single point. So (1) gives a necessary and sufficient condition (in terms of arc lengths) for the chords AD;BE;CF formed by six distinct points A; B; C; D; E; F on a unit circle to meet at a single point. In other words, to give an explicit answer to the question in 6 BJORN POONEN AND MICHAEL RUBINSTEIN the section title, we need to characterize the positive rational solutions to (Here This is a trigonometric diophantine equation in the sense of [2], where it is shown that in theory, there is a finite computation which reduces the solution of such equations to ordinary diophantine equations. The solutions to the analogous equation with only two sines on each side are listed in [9]. If in (2), we substitute multiply both sides by (2i) 3 , and expand, we get a sum of eight terms on the left equalling a similar sum on the right, but two terms on the left cancel with two terms on the right since U leaving \Gammae i-(V +W \GammaU \Gammae i-(Y +Z \GammaX If we move all terms to the left hand side, convert minus signs into e \Gammai- , multiply by we obtainX in which Conversely, given rational numbers necessarily positive) which sum to 1 and satisfy (3), we can recover U; V; W;X;Y;Z, (for example, but we must check that they turn out positive. 3. Zero as a sum of 12 roots of unity In order to enumerate the solutions to (2), we are led, as in the end of the last section, to classify the ways in which 12 roots of unity can sum to zero. More generally, we will study relations of the form a DIAGONALS OF A REGULAR POLYGON 7 where the a i are positive integers, and the j i are distinct roots of unity. (These have been studied previously by Schoenberg [12], Mann [8], Conway and Jones [2], and others.) We call a i the weight of the relation S. (So we shall be particularly interested in relations of weight 12.) We shall say the relation (4) is minimal if it has no nontrivial subrelation; i.e., if implies either b By induction on the weight, any relation can be represented as a sum of minimal relations (but the representation need not be unique). Let us give some examples of minimal relations. For each n - 1, let exp(2-i=n) be the standard primitive n-th root of unity. For each prime p, let R p be the relation Its minimality follows from the irreducibility of the cyclotomic polynomial. Also we can "rotate" any relation by multiplying through by an arbitrary root of unity to obtain a new relation. In fact, Schoenberg [12] proved that every relation (even those with possibly negative coefficients) can be obtained as a linear combination with positive and negative integral coefficients of the R p and their rotations. But we are only allowing positive combinations, so it is not clear that these are enough to generate all relations. In fact it is not even true! In other words, there are other minimal relations. If we subtract R 3 from R 5 , cancel the 1's and incorporate the minus signs into the roots of unity, we obtain a new relation which we will denote (R 5 In general, if S and are relations, we will use the notation to denote any relation obtained by rotating the T i so that each shares exactly one root of unity with S which is different for each i, subtracting them from S, and incorporating the minus signs into the roots of unity. For notational convenience, we will write (R example. Note that although (R 5 unambiguously (up to rotation) the relation listed in (5), in general there will be many relations of type up to rotational equivalence. Let us also remark that including R 2 's in the list of T 's has no effect. It turns out that recursive use of the construction above is enough to generate all minimal relations of weight up to 12. These are listed in Table 1. The completeness and correctness of the table will be proved in Theorem 3 below. Although there are 107 minimal relations up to rotational equivalence, often the minimal relations within Weight Relation type Number of relations of that type 6 (R 5 (R 9 (R (R (R (R (R (R (R (R (R (R (R Table 1. The 107 minimal relations of weight up to 12. one of our classes are Galois conjugates. For example, the two minimal relations of type (R are conjugate under Gal(Q (i 15 )=Q ), as pointed out in [8]. The minimal relations with defined as in (4)) had been previously catalogued in [8], and those with k - 9 in [2]. In fact, the a i in these never exceed 1, so these also have weight less than or equal to 9. Theorem 3. Table 1 is a complete listing of the minimal relations of weight up to 12 (up to rotation). The following three lemmas will be needed in the proof. Lemma 1. If the relation (4) is minimal, then there are distinct primes so that each j i is a -th root of unity, after the relation has been suitably rotated. Proof. This is a corollary of Theorem 1 in [8]. Lemma 2. The only minimal relations (up to rotation) involving only the 2p-th roots of unity, for p prime, are R 2 and R p . DIAGONALS OF A REGULAR POLYGON 9 Proof. Any 2p-th root of unity is of the form \Sigmai i . If both +i i and \Gammai i occurred in the same relation, then R 2 occurs as a subrelation. So the relation has the form By the irreducibility of the cyclotomic polynomial, are independent over Q save for the relation that their sum is zero, so all the c i must be equal. If they are all positive, then R p occurs as a subrelation. If they are all negative, then rotated by -1 (i.e., 180 degrees) occurs as a subrelation. Lemma 3. Suppose S is a minimal relation, and are picked as in Lemma 1 with (or a rotation) is of the form (R ps are minimal relations not equal to R 2 and involving only roots of unity, such that Proof. Since every p 1 -th root of unity is uniquely expressible as a p 1 th root of unity and a p s -th root of unity, the relation can be rewritten as ps where each f i is a sum of p 1 roots of unity, which we will think of as a sum (not just its value). Let Km be the field obtained by adjoining the p 1 roots of unity to Q . the only linear relation satisfied by ps ps over K s\Gamma1 is that their sum is zero. Hence (6) forces the values of the f i to be equal. The total number of roots of unity in all the f i 's is w(S) ! 2p s , so by the pigeonhole principle, some f i is zero or consists of a single root of unity. In the former case, each f j sums to zero, but at least two of these sums contain at least one root of unity, since otherwise s was not minimal, so one of these sums gives a subrelation of S, contradicting its minimality. So some f i consists of a single root of unity. By rotation, we may assume f sums to 1, and if it is not simply the single root of unity 1, the negatives of the roots of unity in f i together with 1 form a relation T i which is not R 2 and involves only roots of unity, and it is clear that S is of type (R ps If one of the T 's were not minimal, then it could be decomposed into two nontrivial subrelations, one of which would not share a root of unity with the R ps , and this would give a nontrivial subrelation of S, contradicting the minimality of S. Finally, w(S) must equal the sum of the weights of R ps and the T 's, minus 2j to account for the roots of unity that are cancelled in the construction of (R ps Proof of Theorem 3. We will content ourselves with proving that every relation of weight up to 12 can be decomposed into a sum of the ones listed in Table 1, it then being straightforward to check that the entries in the table are distinct, and that none of them can be further decomposed into relations higher up in the table. Let S be a minimal relation with w(S) - 12. Pick In particular, p s - 12, so Case 1: Here the only minimal relations are R 2 and R 3 , by Lemma 2. Case 2: If w(S) ! 10, then we may apply Lemma 3 to deduce that S is of type (R Each T must be R 3 (since p by the last equation in Lemma 3. The number of relations of type (R 5 : jR 3 ), up to rotation, is =5. (There are ways to place the R 3 's, but one must divide by 5 to avoid counting rotations of the same relation.) as in (6). If some f i consists of zero or one roots of unity, then the argument of Lemma 3 applies, and S must be of the form (R which contradicts the last equation in the Lemma. Otherwise the numbers of (sixth) roots of unity occurring in f must be 2,2,2,2,2 or 2,2,2,2,3 or 2,2,2,3,3 or 2,2,2,2,4 in some order. So the common value of the f i is a sum of two sixth roots of unity. By rotating by a sixth root of unity, we may assume this value is 0, 1, or 1 If it is 0 or 1, then the arguments in the proof of Lemma 3 apply. So assume it is 1 . The only way two sixth roots of unity can sum to 1+ i 6 is if they are 1 and i 6 in some order. The only ways three sixth roots of unity can sum to 1 are 6 or i 6 6 . So if the numbers of roots of unity occurring in f are 2,2,2,2,2 or 2,2,2,2,3, then S will contain R 5 or its rotation by i 6 , and the same will be true for 2,2,2,3,3 unless the two f i with three terms are 1 6 , in which case S contains (R 5 Finally, it is impossible to as a sum of sixth roots of unity without using 1 or i 6 , so if the numbers are 2,2,2,2,4, then again S contains R 5 or its rotation by . Thus there are no minimal relations S with Case 3: 7, we can apply Lemma 3. Now the sum of w(T required to be w(S) \Gamma 7 which at most 5, so the T 's that may be used are R 3 , R 5 , (R and the two of type (R and 5, respectively. So the problem is reduced to listing the partitions of w(S) \Gamma 7 into parts of size 1, 3, 4, and 5. If all parts used are 1, then we get (R 7, and there are =7 distinct relations in this class. Otherwise exactly one part of size 3, 4, or 5 is used, and the possibilities are as follows. If a part of size 3 is used, we get (R 7 DIAGONALS OF A REGULAR POLYGON 11 Partition Relation type (R (R (R (R (R (R (R (R (R (R Partition Relation type (R R 7 +R 5 (R Table 2. The types of relations of weight 12. (R weights 10, 11, 12 respectively. By rotation, the R 5 may be assumed to share the 1 in the R 7 , and then there are ways to place the R 3 's where i is the number of R 3 's. If a part of size 4 is used, we get (R of weight 11 or (R 7 By rotation, the (R 5 may be assumed to share the 1 in the R 7 , but any of the six roots of unity in the (R 5 may be rotated to be 1. The R 3 can then overlap any of the other 6 seventh roots of unity. Finally, if a part of size 5 is used, we get (R There are two different relations of type (R that may be used, and each has seven roots of unity which may be rotated to be the 1 shared by the R 7 , so there are 14 of these all together. Case 4: Applying Lemma 3 shows that the only possibilities are R 11 of weight 11, and (R Now a general relation of weight 12 is a sum of the minimal ones of weight up to 12, and we can classify them according to the weights of the minimal relations, which form a partition of 12 with no parts of size 1 or 4. We will use the notation (R for example, to denote a sum of three minimal relations of type (R and R 3 . Table 2 lists the possibilities. The parts may be rotated independently, so any category involving more than one minimal relation contains infinitely many relations, even up to rotation (of the entire relation). Also, the categories are not mutually exclusive, because of the non-uniqueness of the decomposition into minimal relations. Figure 3. A surprising trivial solution for the 16-gon. The intersection point does not lie on any of the 16 lines of symmetry of the 16-gon. 4. Solutions to the trigonometric equation Here we use the classification of the previous section to give a complete listing of the solutions to the trigonometric equation (2). There are some obvious solutions to (2), namely those in which U; V; W are arbitary positive rational numbers with sum 1=2, and X; Y; Z are a permutation of U; V; W . We will call these the trivial solutions, even though the three-diagonal intersections they give rise to can look surprising. For example, see Figure 3 for an example on the 16-gon. The twelve roots of unity occurring in (3) are not arbitrary; therefore we must go through Table 2 to see which relations are of the correct form, i.e., expressible as a sum of six roots of unity and their inverses, where the product of the six is -1. Because of the large number of cases, we perform this calculation using Mathematica. Each entry of Table 2 represents a finite number of linearly parameterized (in the exponents) families of relations of weight 12. For each parameterized family, we check to see what additional constraints must be put on the parameters for the relation to be of the form of (3). Next, for each parameterized family of solutions to (3), we calculate the corresponding U; V; W;X;Y;Z and throw away solutions in which some of these are nonpositive. Finally, we sort U; V; W and X; Y; Z and interchange the two triples if U ? X, in order to count the solutions only up to symmetry. The results of this computation are recorded in the following theorem. DIAGONALS OF A REGULAR POLYGON 13 U Table 3. The nontrivial infinite families of solutions to (2). Theorem 4. The positive rational solutions to (2), up to symmetry, can be classified as follows: 1. The trivial solutions, which arise from relations of type 6R 2 . 2. Four one-parameter families of solutions, listed in Table 3, which arise from relations of type 3. Sixty-five "sporadic" solutions, listed in Table 4, which arise from the other types of weight 12 relations listed in Table 2. The only duplications in this list are that the second family of Table 3 gives a trivial solution for 1=12, and that the first and fourth families of Table 3 give the same solution when both. Some explanation of the tables is in order. The last column of Table 3 gives the allowable range for the rational parameter t. The entries of Table 4 are sorted according to the least common denominator of U; V; W;X;Y;Z, which is also the least n for which diagonals of a regular n-gon can create arcs of the corresponding lengths. The reason 11 does not appear in the least common denominator for any sporadic solution is that the relation (R 11 : R 3 ) cannot be put in the form of (3) with the ff j summing to 1, and hence leads to no solutions of (2). (Several other types of relations also give rise to no solutions.) Tables 3 and 4 are the same as Bol's tables at the bottom of page 40 and on page 41 of [1], in a slightly different format. The arcs cut by diagonals of a regular n-gon have lengths which are multiples of 2-=n, so U , V , W , X, Y and Z corresponding to any configuration of three diagonals meeting must be multiples of 1=n. With this additional restriction, trivial solutions to (2) occur only when n is even (and at least 6). Solutions within the infinite families of Table occur when n is a multiple of 6 (and at least 12), and there t must be a multiple of 1=n. Sporadic solutions with least common denominator d occur if and only if n is a multiple of d. 5. Intersections of more than three diagonals Now that we know the configurations of three diagonals meeting, we can check how they overlap to produce configurations of more than three diagonals meeting. 14 BJORN POONEN AND MICHAEL RUBINSTEIN 1=15 1=6 4=15 1=10 1=10 3=10 1=15 1=15 7=15 1=15 1=10 7=30 1=42 3=14 5=14 1=21 1=6 4=21 1=42 1=6 19=42 1=14 2=21 4=21 1=42 1=6 13=42 1=21 1=14 8=21 1=42 1=21 13=21 1=42 1=14 3=14 1=12 2=15 19=60 1=10 3=20 13=60 1=15 11=60 13=60 1=12 1=10 7=20 1=60 4=15 23=60 1=12 1=10 3=20 1=60 4=15 3=10 1=20 1=12 17=60 1=60 13=60 9=20 1=12 1=10 2=15 1=60 13=60 5=12 1=20 2=15 1=6 1=60 1=6 31=60 1=15 1=10 2=15 1=60 1=6 5=12 1=20 1=15 17=60 84 1=12 3=14 19=84 11=84 13=84 4=21 1=14 11=84 23=84 1=12 2=21 29=84 1=42 1=12 7=12 1=21 1=14 4=21 1=84 25=84 5=14 5=84 1=12 4=21 1=84 5=21 5=12 5=84 1=14 17=84 1=84 3=14 37=84 1=21 1=12 17=84 1=84 1=6 43=84 1=21 1=14 4=21 90 1=18 13=90 7=18 11=90 2=15 7=45 1=90 23=90 31=90 2=45 1=15 5=18 1=90 17=90 47=90 1=18 4=45 2=15 1=12 19=120 29=120 1=10 13=120 37=120 1=60 13=120 73=120 1=20 1=12 2=15 1=120 7=20 43=120 7=120 11=120 2=15 1=120 13=60 61=120 1=20 1=12 2=15 1=35 2=15 97=210 1=14 17=210 47=210 Table 4. The solutions to (2). DIAGONALS OF A REGULAR POLYGON 15 We will disregard configurations in which the intersection point is the center of the n-gon, since these are easily described: there are exactly n=2 diagonals (diameters) through the center when n is even, and none otherwise. When k diagonals meet, they form 2k arcs, whose lengths we will measure as a fraction of the whole circumference (so they will be multiples of 1=n) and list in counterclockwise order. (Warning: this is different from the order used in Tables 3 and 4.) The least common denominator of the numbers in this list will be called the denominator of the configuration. It is the least n for which the configuration can be realized as diagonals of a regular n-gon. Lemma 4. If a configuration of k - 2 diagonals meeting at an interior point other than the center has denominator dividing d, then any configuration of diagonals meeting at that point has denominator dividing LCM(2d; 3). Proof. We may assume 2. Any other configuration of diagonals through the intersection point is contained in the union of configurations obtained by adding one diagonal to the original two, so we may assume the final configuration consists of three diagonals, two of which were the original two. Now we need only go through our list of three-diagonal intersections. It can be checked (using Mathematica) that removing any diagonal from a sporadic configuration of three intersecting diagonals yields a configuration whose denominator is the same or half as much, except that it is possible that removing a diagonal from a three-diagonal configuration of denominator 210, yields one of denominator 70, which proves the desired result for this case. The additive group generated by 1=6 and the normalized arc lengths of a configuration obtained by removing a diagonal from a configuration corresponding to one of the families of Table 3 contains 2t where t is the parameter, (as can be verified using Mathematica again), which means that adding that third diagonal can at most double the denominator (and throw in a factor of 3, if it isn't already there). Similarly, it is easily checked (even by hand), that the subgroup generated by the normalized arc lengths of a configuration obtained by removing one of the three diagonals of a configuration corresponding to a trivial solution to (2) but with intersection point not the center, contains twice the arc lengths of the original configuration. Corollary 1. If a configuration of three or more diagonals meeting includes three forming a sporadic configuration, then its denominator is 30, 42, 60, 84, 90, 120, 168, 180, 210, 240, or 420. Proof. Combine the lemma with the list of denominators of sporadic configurations listed in Table 4. For k - 4, a list of 2k positive rational numbers summing to 1 arises this way if and only if the lists of length 2k \Gamma 2 which would arise by removing the first or second diagonal actually correspond to intersecting diagonals. Suppose 4. If we Range Table 5. The one-parameter families of four-diagonal configurations. specify the sporadic configuration or parameterized family of configurations that arise when we remove the first or second diagonal, we get a set of linear conditions on the eight arc lengths. Corollary 1 tells us that we get a configuration with denominator among 30, 42, 60, 84, 90, 120, 168, 180, 210, 240, and 420, if one of these two is sporadic. Using Mathematica to perform this computation for the rest of possibilities in Theorem 4 shows that the other four-diagonal configurations, up to rotation and reflection, fall into 12 one-parameter families, which are listed in Table 5 by the eight normalized arc lengths and the range for the parameter t, with a finite number of exceptions of denominators among 12, 18, 24, 30, 36, 42, 48, 60, 84, and 120. We will use a similar argument when 5. Any five-diagonal configuration containing a sporadic three-diagonal configuration will again have denominator among 30, 42, 60, 84, 90, 120, 168, 180, 210, 240, and 420, again. Any other five-diagonal configuration containing one of the exceptional four-diagonal configurations will have denominator among 12, 18, 24, 30, 36, 42, 48, 60, 72, 84, 96, 120, 168, and 240, by Lemma 4. Finally, another Mathematica computation shows that the one-parameter families of four-diagonal configurations overlap to produce the one-parameter families listed (up to rotation and reflection) in Table 6, and a finite number of exceptions of denominators among 12, 18, 24, and 30. For six-diagonal configuration containing a sporadic three- diagonal configuration will again have denominator among 30, 42, 60, 84, 90, 120, 168, 180, 210, 240, and 420. Any six-diagonal configuration containing one of the exceptional four-diagonal configurations will have denominator among 12, 18, 24, 30, 36, 42, 48, 60, 72, 84, 96, 120, 168, and 240. Any six-diagonal configuration containing one of the exceptional five-diagonal configurations will have denominator among 12, 18, 24, 30, 36, 48, and 60. Another Mathematica computation shows that DIAGONALS OF A REGULAR POLYGON 17 Range Table 6. The one-parameter families of five-diagonal configurations. the one-parameter families of five-diagonal configurations cannot combine to give a six-diagonal configuration. Finally for k - 7, any k-diagonal configuration must contain an exceptional configuration of 3, 4, or 5 diagonals, and hence by Lemma 4 has denominator among 12, We summarize the results of this section in the following. Proposition 1. The configurations of k - 4 diagonals meeting at a point not the center, up to rotation and reflection, fall into the one-parameter families listed in Tables 5 and 6, with finitely many exceptions (for fixed of denominators among 12, 18, 24, 30, 36, 42, 48, 60, 72, 84, 90, 96, 120, 168, 180, 210, 240, and 420. In fact, many of the numbers listed in the proposition do not actually occur as denominators of exceptional configurations. For example, it will turn out that the only denominator greater than 120 that occurs is 210. 6. The formula for intersection points Let a k (n) denote the number of points inside the regular n-gon other than the center where exactly k lines meet. Let b k (n) denote the number of k-tuples of diagonals which meet at a point inside the n-gon other than the center. Each interior point at which exactly m diagonals meet gives rise to such k-tuples, so we have the relationship m-k am (n) (7) Since every four distinct vertices of the n-gon determine one pair of diagonals which intersect inside, the number of such pairs is exactly , but if n is even, then of these are pairs which meet at the center, so (Recall that ffi m (n) is defined to be 1 if n is a multiple of m, and 0 otherwise.) We will use the results of the previous two sections to deduce the form of b k (n) and then the form of a k (n). To avoid having to repeat the following, let us make a definition. Definition . A function on integers n - 3 will be called tame if it is a linear combination (with rational coefficients) of the functions n 3 , Proposition 2. For each k - 2, the function b k (n)=n on integers n - 3 is tame. Proof. The case handled by (8), so assume k - 3. Each list of 2k normalized arc lengths as in Section 5 corresponding to a configuration of k diagonals meeting at a point other than the center, considered up to rotation (but not reflection), contributes n to b k (n). (There are n places to start measuring the arcs from, and these n configurations are distinct, because the corresponding intersection points differ by rotations of multiples of 2-=n, and by assumption they are not at the center.) counts such lists. 3. When n is even, the family of trivial solutions to the trigonometric equation (2) has are positive integers with sum n=2, and X, Y , and Z are some permutation of U , V , W . Each permutation gives rise to a two-parameter family of six-long lists of arc lengths, and the number of lists with each family is the number of partitions of n=2 into three positive parts, which is a quadratic polynomial in n. Similarly each family of solutions in Table 3 gives rise to a number of one-parameter families of lists, when n is a multiple of 6, each containing dn=6e \Gamma 1 or dn=12e \Gamma 1 lists. These functions of n (extended to be when 6 does not divide n) are expressible as a linear combination of nffi 6 (n), ffi 6 (n), and Finally the sporadic solutions to 2 give rise to a finite number of lists, having denominators among 30, 42, 60, 84, 90, 120, and 210, so their contribution to 3 (n)=n is a linear combination of ffi But these families of lists overlap, so we must use the Principle of Inclusion-Exclusion to count them properly. To show that the result is a tame function, it suffices to show that the number of lists in any intersection of these families is a tame function. When two of the trivial families overlap but do not coincide, they overlap where two of the a, b, and c above are equal, and the corresponding lists lie in one of the one-parameter families (t; t; t; t; (with 1=4), each of which contain dn=4e \Gamma 1 lists (for n even). This function of n is a combination of nffi 2 (n), hence it is tame. Any other intersection of the infinite families must contain the intersection of two one-parameter families which are among the two above or arise from Table 3, and a Mathematica computation shows that such an intersection consists of at most a single list of denominator among 6, 12, 18, 24, and 30. And, of course, any intersection involving a single sporadic list, can contain at most that sporadic list. Thus the number of lists DIAGONALS OF A REGULAR POLYGON 19 within any intersection is a tame function of n. Finally we must delete the lists which correspond to configurations of diagonals meeting at the center. These are the lists within the trivial two-parameter family (t; u; number is also a tame function of n, by the Principle of Inclusion-Exclusion again. Thus b 3 (n)=n is tame. Next suppose 4. The number of lists within each family listed in Table 5, or the reflection of such a family, is (when n is divisible by 6) the number of multiples of 1=n strictly between ff and fi, where the range for the parameter t is ff This number is dfine the table shows that ff and fi are always multiples of 1=24, this function of n is expressible as a combination of nffi 6 (n) and a function on multiples of 6 depending only on n mod 24, and the latter can be written as a combination of ffi 6 (n), so it is tame. Mathematica shows that when two of these families are not the same, they intersect in at most a single list of denominator among 6, 12, 18, and 24. So these and the exceptions of Proposition 1 can be counted by a tame function. Thus, again by the Principle of Inclusion-Exclusion, b 4 (n)=n is tame. The proof for identical to that of 4, using Table 6 instead of Table 5, and using another Mathematica computation which shows that the intersections of two one-parameter families of lists consist of at most a single list of denominator 24. The proof for k - 6 is even simpler, because then there are only the exceptional lists. By Proposition 1, b k (n)=n is a linear combination of ffi m (n) where m ranges over the possible denominators of exceptional lists listed in the proposition, so it is tame. Lemma 5. A tame function is determined by its values at 10, 12, 18, 24, 30, 36, 42, 48, 54, 60, 66, 72, 84, 90, 96, 120, 168, 180, 210, and 420. Proof. By linearity, it suffices to show that if a tame function f is zero at those values, then f the zero linear combination of the functions in the definition of a tame function. The vanishing at forces the coefficients of n 3 , 1 to vanish, by Lagrange interpolation. Then comparing the values at shows that the coefficient of ffi 4 (n) is zero. The vanishing at forces the coefficients of to vanish. Comparing the values at shows that the coefficient of nffi 6 is zero. Comparing the values at shows that the coefficient of At this point, we know that f(n) is a combination of ffi m (n), for 30, 36, 42, 48, 60, 72, 84, 90, 96, 120, 168, 180, 210, and 420. For each m in turn, now implies that the coefficient of ffi m (n) is zero. Proof of Theorem 1. Computation (see the appendix) shows that the tame function b 8 (n)=n vanishes at all the numbers listed in Lemma 5. Hence by that lemma, b 8 0 for all n. Thus by (7), a k (n) and b k (n) are identically zero for all k - 8 as well. By reverse induction on k, we can invert (7) to express a k (n) as a linear combination of b m (n) with m - k. Hence a k (n)=n is tame as well for each k - 2. Computation shows that the equations a a 3 a 4 a 5 a 6 a 7 hold for all the n listed in Lemma 5, so the lemma implies that they hold for all - 3. These formulas imply the remarks in the introduction about the maximum number of diagonals meeting at an interior point other than the center. Finally a k (n) a k (n); which gives the desired formula. (The in the expression for I(n) is to account for the center point when n is even, which is the only point not counted by the a k .) 7. The formula for regions We now use the knowledge obtained in the proof of Theorem 1 about the number of interior points through which exactly k diagonals pass to calculate the number of regions formed by the diagonals. Proof of Theorem 2. Consider the graph formed from the configuration of a regular n-gon with its diagonals, in which the vertices are the vertices of the n-gon together with the interior intersection points, and the edges are the sides of the n-gon together DIAGONALS OF A REGULAR POLYGON 21 with the segments that the diagonals cut themselves into. As usual, let V denote the number of vertices of the graph, E the number of edges, and F the number of regions formed, including the region outside the n-gon. We will employ Euler Formula's 2. I(n). We will count edges by counting their ends, which are 2E in number. Each vertex has the center (if n is even) has n edge ends, and any other interior point through which exactly k diagonals pass has 2k edge ends, so So the desired number of regions, not counting the region outside the n-gon, is Substitution of the formulas derived in the proof of Theorem 1 for a k (n) and I(n) yields the desired result. Appendix: computations and tables In Table 7 we list I(n); R(n); a To determine the polynomials listed in Theorem 1 more data was needed especially for The largest n for which this was required was 420. For speed and memory conser- vation, we took advantage of the regular n-gon's rotational symmetry and focused our attention on only 2-=n radians of the n-gon. The data from this computation is found in Table 8. Although we only needed to know the values at those n listed in Lemma 5 of Section 6, we give a list for so that the nice patterns can be seen. The numbers in these tables were found by numerically computing (using a C program and 64 bit precision) all possible intersections, and sorting them by their x coordinate. We then focused on runs of points with close x coordinates, looking for points with close y coordinates. Several checks were made to eliminate any fears (arising from round-off errors) of distinct points being mistaken as close. First, the C program sent data to Maple which checked that the coordinates of close points agreed to at least 40 decimal places. Second, we verified for each n that close points came in counts of the form diagonals meeting at a point give rise to close points. Hence, any run whose length is not of this form indicates a computational error). A second program was then written and run on a second machine to make the computations completely rigorous. It also found the intersection points numerically, sorted them and looked for close points, but, to be absolutely sure that a pair of close 22 BJORN POONEN AND MICHAEL RUBINSTEIN points were actually the same, it checked that for the two pairs of diagonals respectively, the triples l 1 each divided the circle into arcs of lengths consistent with Theorem 4. Since this test only involves comparing rational numbers, it could be performed exactly. A word should also be said concerning limiting the search to 2-=n radians of the n-gon. Both programs looked at slightly smaller slices of the n-gon to avoid problems caused by points near the boundary. More precisely, we limited our search to points whose angle with the origin fell between [c 1 of made sure not to include the origin in the count. Here " was chosen to be :00000000001 and c 1 was chosen to be :00000123 would have led to problems since there are many intersection points with angle 0 or 2-=n). To make sure that no intersection points were omitted, the number of points found (counting multiplicity) was compared with ( Acknowledgements We thank Joel Spencer and Noga Alon for helpful conversations. Also we thank Jerry Alexanderson, Jeff Lagarias, Hendrik Lenstra, and Gerry Myerson for pointing out to us many of the references below. --R Beantwoording van prijsvraag No. Trigonometric Diophantine equations (On vanishing sums of roots of unity) On a problem of Steinhaus about polygons Diagonalen im regularen n-Eck Number of intersections of diagonals in regular n-gons On linear relations between roots of unity Rational products of sines of rational angles Adventitious quadrangles: a geometrical approach Multiple intersections of diagonals of regular polygons A note on the cyclotomic polynomial Mathematical Snapshots Problem 225 Adventitious angles --TR

intersection points;diagonals;adventitious quadrangles;regular polygons;roots of unity

288868

Rankings of Graphs.

A vertex (edge) coloring $\phi:V\rightarrow \{1,2,\ldots ,t\}$ ($\phi':E\rightarrow \{1,2,\ldots,$ $t\}$) of a graph G=(V,E) is a vertex (edge) t-ranking if, for any two vertices (edges) of the same color, every path between them contains a vertex (edge) of larger color. The {\em vertex ranking number} $\chi_{r}(G)$ ({\em edge ranking number} $\chi_{r}'(G)$) is the smallest value of t such that G has a vertex (edge) t-ranking. In this paper we study the algorithmic complexity of the {\sc Vertex Ranking} and {\sc Edge Ranking} problems. It is shown that $\chi_{r}(G)$ can be computed in polynomial time when restricted to graphs with treewidth at most k for any fixed k. We characterize the graphs where the vertex ranking number $\chi_{r}$ and the chromatic number $\chi$ coincide on all induced subgraphs, show that $\chi_{r}(G)=\chi (G)$ implies $\chi (G)=\omega (G)$ (largest clique size), and give a formula for $\chi_{r}'(K_n)$.

Introduction In this paper we consider vertex rankings and edge rankings of graphs. The vertex ranking problem, also called the ordered coloring problem [15], has received much attention lately because of the growing number of applications. There are applications in scheduling problems of assembly steps in manufacturing systems [19], e.g., edge ranking of trees can be used to model the parallel assembly of a product from its components in a quite natural manner [6, 12, 13, 14]. Department of Computer Science, Utrecht University, P.O. Box 80.089, 3508 the Netherlands. Email: [email protected]. This author was partially supported by the ESPRIT Basic Research Actions of the EC under contract 7141 (project ALCOM II) y Department of Computer Science and Engineering, University of Nebraska - Lincoln, Lincoln, NE 68588-0115, U.S.A. This author was partially supported by the Office of Naval Research under Grant No. N0014-91-J-1693 z Institut f?r Informatik, TU M-unchen, 80290 M-unchen, Germany. On leave from Univer- sit-at Trier, Germany x Department of Mathematics and Computing Science, Eindhoven University of Technology, P.O.Box 513, 5600 MB Eindhoven, the Netherlands - IRISA, Campus Universitaire de Beaulieu, 35042 Rennes Cedex, France. On leave from Friedrich-Schiller-Universit?t, Jena. This author was partially supported by Deutsche under Kr 1371/1-1 k Fakult?t f?r Mathematik und Informatik, Friedrich-Schiller-Universit?t, Universit?ts- hochhaus, 07740 Jena, Germany Computer and Automation Institute, Hungarian Academy of Sciences, H-1111 Budapest, Kende u. 13-17, Hungary. This author was partially supported by the "OTKA" Research Fund of the Hungarian Academy of Sciences, Grant No. 2569 Furthermore the problem of finding an optimal vertex ranking is equivalent to the problem of finding a minimum-height elimination tree of a graph [6, 7]. This measure is of importance for the parallel Cholesky factorization of matrices [3, 9, 18]. Yet other applications lie in the field of VLSI-layout [17, 26]. The vertex ranking problem 'Given a graph G and a positive integer t, decide whether - r (G) - t ' is NP-complete even when restricted to cobipartite graphs since Pothen has shown that the equivalent minimum elimination tree height problem remains NP-complete on cobipartite graphs [20]. A short proof of the NP-completeness of vertex ranking is given in Section 3. Much work has been done in finding optimal rankings of trees. For trees there is a linear-time algorithm finding an optimal vertex ranking [24]. For the closely related edge ranking problem on trees a O(n 3 ) algorithm was given in [8]. Recently, Zhou and Nishizeki obtained an O(n log n) algorithm for optimally edge ranking trees [28] (see also [29]). Efficient vertex ranking algorithms for permutation, trapezoid, interval, circular-arc, circular permutation graphs, and cocompara- bility graphs of bounded dimension are presented in [7]. Moreover, the vertex ranking problem is trivial on split graphs and it is solvable in linear time on cographs [25]. In [15], typical graph theoretical questions, as they are known from the coloring theory of graphs, are investigated. This also leads to a O( n) bound for the vertex ranking number of a planar graph and the authors describe a polynomial-time algorithm which finds a vertex ranking of a planar graph using only O( n) colors. For graphs in general there is an approximation algorithm of performance ratio O(log 2 n) for the vertex ranking number [3, 16]. In [3] it is also shown that one plus the pathwidth of a graph is a lower bound for the vertex ranking number of the graph (hence a planar graph has pathwidth O( n), which is also shown in [16] using different methods). Our goal is to extend the known results in both the algorithmic and graph theoretic directions. The paper is organized as follows. In Section 2 the necessary notions and preliminary results are given. We study the algorithmic complexity of determining whether a graph G fulfills - r (G) - t and - 0 respectively, in Sections 3, 4, and 5. In Section 6 we characterize those graphs for which the vertex ranking number and the chromatic number coincide on every induced subgraph. Those graphs turn out to be precisely those containing no path and cycle on four vertices as an induced subgraph; hence, we obtain a characterization of the trivially perfect graphs [11] in terms of rankings. Moreover we show that implies that the chromatic number of G is equal to its largest clique size. In Section 7 we give a recurrence relation allowing us to compute the edge ranking number of a complete graph. Preliminaries We consider only finite, undirected and simple graphs E). Throughout the paper n denotes the cardinality of the vertex set V and m denotes that of the edge set E of the graph E). For graph-theoretic concepts, definitions and properties of graph classes not given here we refer to [4, 5, 11]. E) be a graph. A subset U ' V is independent if each pair of vertices nonadjacent. A graph E) is bipartite if there is a partition of V into two independent sets A and B. The complement of the graph E) is the graph G having vertex set V and edge set ffv; wg j v 6= Eg. For W ' V we denote by G[W ] the subgraph of E) induced by the vertices of W , and for X ' E we write G[X] for the graph and edge set X. E) be a graph and let t be a positive integer. A (vertex) t-ranking, called ranking for short if there is no ambiguity, is a coloring such that for every pair of vertices x and y with and for every path between x and y there is a vertex z on this path with c(z) ? c(x). The vertex ranking number of G, - r (G), is the smallest value t for which the graph G admits a t-ranking. By definition adjacent vertices have different colors in any t-ranking, thus any t- ranking is a proper t-coloring. Hence - r (G) is bounded below by the chromatic number -(G). A vertex - r (G)-ranking of G is said to be an optimal (vertex) ranking of G. The edge ranking problem is closely related to the vertex ranking problem. E) be a graph and let t be a positive integer. An edge t-ranking is an edge coloring c 0 such that for every pair of edges e and f with c 0 and for every path between e and f there is an edge g on this path with c 0 (g) ? c 0 (e). The edge ranking number - 0 r (G) is the smallest value of t such that G has an edge t-ranking. Remark 3 There is a one-to-one correspondence between the edge t-rankings of a graph G and the vertex t-rankings of its line graph L(G). Hence - 0 An edge t-ranking of a graph G is a particular proper edge coloring of G. Hence r (G) is bounded below by the chromatic index - 0 (G). An edge - 0 r (G)-ranking of G is said to be an optimal edge ranking of G. As shown in [7], the vertex ranking number of a connected graph is equal to its minimum elimination tree height plus one. Thus (vertex) separators and edge separators are a convenient tool for investigating rankings of graphs. A subset of a graph E) is said to be a separator if G[V n S] is disconnected. A subset R ' E of a graph E) is said to be an edge separator (or edge cut) if G[E n R] is disconnected. In this paper we use the separator tree for studying vertex rankings. This concept is closely related to elimination trees (cf.[3, 7, 18]). Definition 4 Given a vertex t-ranking of a connected graph E), we assign a rooted tree T (c) to it by an inductive construction, such that a separator of a certain induced subgraph of G is assigned to each internal node of T (c) and the vertices of each set assigned to a leaf of T (c) have different colors: 1. If no color occurs more than once in G, then T (c) consists of a single vertex r (called root), assigned to the vertex set of G. 2. Otherwise, let i be the largest color assigned to more than one vertex by c. Then has to be a separator S of G. We create a root r of T (c) and assign S to r. (The induced subgraph of G corresponding to the subtree of T rooted at r will be G itself.) Assuming that a separator tree has already been defined for each connected component G i of the graph G[V n S], the children of r in T (c) will be the vertices r i and the subtree of T (c) rooted at r i will be T i (c). The rooted tree T (c) is said to be a separator tree of G. Notice that all vertices of G assigned to nodes of T (c) on a path from a leaf to the root have different colors. 3 Unbounded ranking It is still unknown whether the edge ranking problem 'Given a graph G and a positive integer t, decide whether - 0 Remark 3 this problem is equivalent to the vertex ranking problem 'Given a graph G and a positive integer t, decide whether - r (G) - t ' when restricted to line graphs. On the other hand, it is a consequence of the NP-completeness of the minimum elimination tree height problem shown by Pothen in [20] and the equivalence of this problem with the vertex ranking problem [6, 7] that the latter is NP-complete even when restricted to graphs that are the complement of bipartite graphs, the so-called cobipartite graphs. For reasons of self-containedness, we start with a short proof of the NP-completeness of vertex ranking, when restricted to cobipartite graphs. The following problem, called balanced complete bipartite subgraph (abbre- viated bcbs) is NP-complete. This is problem [GT24] of [10]. Instance: A bipartite graph E) and a positive integer k. Question: Are there two disjoint subsets such that and such that u implies that Theorem 5 vertex ranking is NP-complete for a cobipartite graphs. Proof: Clearly the problem is in NP. NP-hardness is shown by reduction from bcbs. Let a bipartite graph E) and a positive integer k be given. Let G be the complement of G, thus G is a cobipartite graph. We claim that G has a balanced complete bipartite subgraph with 2 \Delta k vertices if and only if G has a (n \Gamma k)-ranking. Suppose we have sets W 1 and such that for all u We now construct a k- ranking of G. Write W n\Gamma2\Deltak g. We define a vertex ranking c of G as follows: c(v (1) One easily observes that c is a vertex (n \Gamma k)-ranking. Next, let c be a (n\Gammak)-ranking for G. Since G is a cobipartite graph, for each color, there can be at most two vertices with that color, one lying in V 1 and the other in V 2 . Therefore, we have k pairs v (1) with c(v (1) can assume that W Now we show that the subgraph induced by the set W 1 [W 2 forms a balanced complete bipartite subgraph in G. To show this, we prove that each pair of vertices not adjacent in G. Suppose v (1) are adjacent in G. Then, the colors of these vertices must be different. Furthermore, assume w.l.o.g., that c(v (1) we have a path (v (1) c(v (1) contradicting the fact that c is a ranking. Hence the subgraph induced by W 1 [W 2 is indeed a balanced complete bipartite subgraph. This proves the claim, and the NP-completeness of vertex ranking. 2 We show that the analogous result holds for bipartite graphs as well. Theorem 6 vertex ranking remains NP-complete for bipartite graphs. Proof: The transformation is from vertex ranking for arbitrary graphs without isolated vertices. Given the graph G, we construct a graph G and Clearly, the constructed graph G 0 is a bipartite graph. Now we show that G has a t-ranking if and only if G 0 has a (t 1)-ranking. Suppose G has a t-ranking tg. We construct a coloring - c for G 0 in the following way. For the vertices the vertices c is a 1)-ranking of G 0 . On the other hand, let +1g be a (t +1)-ranking of G 0 . We show that - c(v) ? 1 for every vertex . Suppose not and let v be a vertex of V with - wg be an edge incident to v in G. Hence v is adjacent to (e; 1); 1)-ranking, there are l; l 0 with l 6= l 0 such that - implying a path (e; the assumption that - c is a ranking. This proves that - c(v) ? 1 holds for every vertex As a consequence, for each edge there is a vertex to - c((e; 1)-ranking of G 0 . Now we define . The coloring c is a t-ranking of G since the existence of a path between two vertices v and w of G such that c(w) and all inner vertices have smaller colors implies the existence of a path from v to w in G 0 with - c(w) and all inner vertices having smaller colors, contradicting the fact that - c is a 1)-ranking of G 0 . 2 4 Bounded ranking We show that the 'bounded' ranking problems 'Given a graph G, decide whether are solvable in linear time for any fixed t. This will be done by verifying that the corresponding graph classes are closed under certain operations. Definition 7 An edge contraction is an operation on a graph G replacing two adjacent vertices u and v of G by a vertex adjacent to all vertices that were adjacent to u or v. An edge lift is an operation on a graph G replacing two adjacent edges fv; wg and fu; wg of G by one edge fu; vg. Definition 8 A graph H is a minor of the graph G if H can be obtained from G by a series of the following operations: vertex deletion, edge deletion, and edge contraction. A graph class G is minor closed if every minor H of every graph G 2 G also belongs to G. Lemma 9 The class of graphs satisfying - r (G) - t is minor closed for any fixed t. Proof: Since vertex/edge deletion cannot create new paths between monochromatic pairs of vertices, we only have to show that edge contraction does not increase the ranking number. Let E) be a graph with t, and assume is obtained from G by contracting the edge new vertex c uv. Suppose c is a t-ranking of G. We construct of H as follows. uv Suppose - c is not a t-ranking of H. Then there is a path c is a t-ranking of G the vertex c uv must occur in the path. Depending on its neighbors in P we can 'decontract' c uv in the path P into u, v, getting a path P 0 of G violating the ranking condition, in contradiction to the choice of c. 2 Corollary 10 For each fixed t, the class of graphs satisfying - r (G) - t can be recognized in linear time. Proof: In [1], using results from Robertson and Seymour [22, 23], it is shown that every minor closed class of graphs that does not contain all planar graphs, has a linear time recognition algorithm. The result now follows directly from Lemma 9. 2 As regards edge rankings, a simple argument yields a much stronger assertion as follows. Theorem 11 For each fixed t, the class of connected graphs satisfying - 0 t can be recognized in constant time. Proof: For any fixed t, there are only a finite number of connected graphs G with t, as necessary conditions are that the maximum degree of G is at most t, and the diameter of G is bounded by 2 t \Gamma 1. 2 Certainly, the above theorem immediately implies that the graphs G satisfying can be recognized in linear time, by inspecting the connected components separately. This result might have also been obtained via more involved methods, by using results of Robertson and Seymour on graph immersions [21]. Similarly, one can show that for fixed t and d, the class of connected graphs with - r (G) - t and maximum vertex degree d can be recognized in constant time. Definition 12 A graph H is an immersion of the graph G if H can be obtained from G by a series of the following operations: vertex deletion, edge deletion and edge lift. A graph class G is immersion closed if every immersion H of a graph G 2 G also belongs to G. The proof of the following lemma is similar to the one of Lemma 9 and therefore omitted. Lemma 13 The class of graphs satisfying - 0 closed for any fixed t. Linear-time recognizability of the class of graphs satisfying - 0 now also follows from Lemma 13, the results of Robertson and Seymour, and the fact that graphs with - 0 have treewidth at most 2t 2. 5 Computing the vertex ranking number on graphs with bounded treewidth In this section, we show that one can compute - r (G) of a graph G with treewidth at most k in polynomial time, for any fixed k. Such a graph is also called a partial k-tree. This result implies polynomial time computability of the vertex ranking number for any class of graphs with a uniform upper bound on the treewidth, e.g., outerplanar graphs, series-parallel graphs, Halin graphs. The notion of treewidth has been introduced by Robertson and Seymour (see e.g., [22]). Definition 14 A tree-decomposition of a graph E) is a pair (fX a collection of subsets of V , and tree, such that ffl for all edges fv; wg 2 E there is an i 2 I with v; w ffl for all is on the path from i to k in T , then The width of a tree-decomposition (fX 1. The treewidth of a graph E) is the minimum width over all tree-decompositions of G. We often abbreviate (fX When the treewidth of E) is bounded by a constant k, one can find in O(n) time a tree-decomposition (X; T ) of width at most k, such that I = O(n) and T is a rooted binary tree [1]. Denote the root of T as r. We say (X; T ) is a rooted binary tree-decomposition. Definition 15 A terminal graph is a triple (V; E; Z), with (V; E) an undirected graph, and Z ' V a subset of the vertices, called the terminals. To each node i of a rooted binary tree-decomposition (X; T ) of graph E), we associate the terminal graph G is a descendant of ig, and g. As shorthand notation we write p(v; w; G; c; ff), iff there is a path in G from v to w with all internal vertices having colors, smaller than ff under coloring c. If G; c; ff), we denote with P (v; w; G; c; ff) the set of paths in G from v to w with all internal vertices having colors (using color function c), smaller than ff. In the following, suppose t is given. be a terminal graph, and let c be a vertex t-ranking of (V; E). The characteristic of c, Y (c), is the quadruple (cj Z ffl cj Z is the function c, restricted to domain Z. defined true if and only or there is a vertex x 2 V with G; c; i). defined and only if there is a vertex x 2 V with G; c; i) and G; c; i). t; 1g is defined by: f 3 (v; w) is the smallest integer t 0 such that p(v; w; G; c; t 0 ). If there is no path from v to w in G, then Definition 17 A set of characteristics S of vertex t-rankings of a terminal graph G is a full set of characteristics of vertex t-rankings for G (in short: a full set for G), if and only if for every vertex t-ranking c of G, Y (c) 2 S. set C of vertex t-rankings of a terminal graph G is an example set of vertex t-rankings for G (in short: an example set for G), if and only if for every vertex t-ranking c of G, there is an c 0 2 C with Y the set of characteristics of the elements of C forms a full set of characteristics of vertex t-rankings for G. then a full set of characteristics of vertex t-rankings of (with jZj polynomial in V : there are O(log k+1 n) possible values for cj Z , 2 O((k+1) log n) possible values for f 1 , log n) possible values for f 2 , and there are O(log 1k(k+1) n) possible values for f 3 . The following lemma, given in [3], shows that we can ensure this property for graphs with treewidth at most k for fixed k. Lemma If the treewidth of E) is at most k, then - r log n). Let (X; T ) be a rooted binary tree-decomposition of G. Suppose j 2 I is a descendant of i 2 I in T . Suppose c is a vertex t-ranking of G i . The restriction of c to G j is the function cj G j defined by 8v Clearly, cj G j is a vertex t-ranking of G j . If c 0 is another vertex t-ranking of G j , we define the function R(c; c Lemma 19 Let (X; T ) be a rooted binary tree-decomposition of E). Let j be a descendant of i. Let c be a vertex t-ranking of G i , and c 0 be a vertex t-ranking of G j . If Y (cj G j vertex t-ranking of G i , and Y Proof: For brevity, we write c and we write Y (cj G j We start with proving two claims. Proof: Let v; w 2 W 1 , and suppose we have a path We consider those parts of the path p that are part of G such that each p ff (0 - ff - r) is a path with all vertices in W 1 , and each p 0 a path in G j . (Each path is a collection of successive edges, i.e., the last vertex of a path is the first vertex of the next path.) Write v ff for the first vertex on path ff and w ff for the last vertex on path p 0 1). Note that . We now have that there also exists a path p 00 (In words: there exists a path from v to w in G j such that all colors of internal vertices are smaller than t 0 , using coloring c (or, equivalently cj G j As cj G j and c 0 have the same characteristics, there also exists such a path using color function c 0 .) Now, the path formed by the sequence (p is a path in G i between v and w with all colors of internal vertices smaller than t 0 , hence p(v; w; G shows: p(v; w; G can be shown in the same way. 2 there exists a vertex only if there exists a vertex w (w claim 20, we have p(v; w; G x be the last vertex on a path that belongs to W 1 . Write is the last vertex of p 0 and the first vertex of p 00 . p exists a path q Using equality of the characteristics of cj G j and c 0 , we have that there exists a vertex a path from v to w 0 in G i with all internal vertices of color (under color function c 00 ) smaller that t 0 , hence p(v; w The reverse implication of the claim can be shown in a similar way. 2 We now show that c 00 is a vertex t-ranking, or, equivalently, that for all We consider four cases: 1. v; w c is not a vertex ranking, contradiction. 2. there exists a again c is not a vertex ranking, contradiction. 3. w . Similar to Case 2. 4. v; w all vertices on p belong to W 2 , then p is a path in G j , and hence c 0 was not a vertex ranking of G j , contradiction. So, there exist vertices on p that belong to W 1 . Let x be the first vertex on p that belongs to W 1 . Then must exist vertices with a path from v 0 to w 0 with all internal vertices of color (with color function c) less than t 0 . Hence c is not a vertex ranking, contradiction. It remains to show that Y . Suppose 3 ). It follows directly from Consider v; w be the vertex with true. If x can and y 2 X j . (Let y be the last vertex in X j on p 1 .) Similarly, we can write This implies that f 2 (y; z; t 0 ) is true. Hence, there is a vertex x 0 with paths Also, by Lemma 20, we have paths p 0 using path (p 0 12 ) from v to x 0 and path (p 0 22 ) from w to x 0 , it follows that almost identical argument shows . Finally, it follows directly from Claim 20 that g We now describe our algorithm. After a rooted binary tree-decomposition E) has been found (in linear time [1]), the algorithm computes a full set and an example set for every node i 2 I, in a bottom-up order. Clearly, when we have a full set for the root node of T , we can determine whether G has a vertex t-ranking, as we only have to check whether the full set of the root is non-empty. If so, any element of the example set of the root node gives us a vertex t-ranking of G. It remains to show that we can compute for any node i 2 I a full set and an example set, given a full set and an example set for each of the children of i 2 I. This is straightforward for the case that i is a leaf node: enumerate all functions c for each such function c, test whether it is a vertex t-ranking of G i , and if so, put c in the example set, and Y (c) in the full set of characteristics. Next suppose i 2 I has two children j 1 and j 2 . (If i has one child we can add another child j 2 , which is a leaf in T and has we have example sets . We compute a full set S and an example set Q for G i in the following way: Initially, we take S and Q to be empty. For each triple is an element of Q 1 , c 2 is an element of 3 is an arbitrary function c do the ffl Check whether for all (v). If not, skip the following steps and proceed with the next triple. ffl Compute the function c : defined as follows: ffl Check whether c is a vertex t-ranking of G i . If not, skip the following steps and proceed with the next triple. ffl If Y (c) 62 S, then put Y (c) in S and put c in Q. We claim that the resulting sets S and Q form a full set and an example set for G i . Consider an arbitrary vertex t-ranking c 0 of G i . Let c 1 2 Q 1 be the vertex t-ranking of Y that has the same characteristic as c 0 . By definition of example set, c 1 must exist. Similarly, let c be defined by c 3 ). When the algorithm processes the triple first test will hold. Suppose c is the function, computed in the second test. Now note that Hence, by Lemma 19, c is a vertex t-ranking and has the same characteristic as c 0 . Hence, S will contain Y (c), and Q will contain a vertex t-ranking of G i with the same characteristic as c and c 0 . As the size of a full set, and hence of an example set for graphs G i , i 2 I is polynomial, it follows that the computation of a full set and example set from these sets associated with the children of the node, can be done in polynomial time. (There are a polynomial number of triples For each triple, the computation given above costs polynomial time.) As there are a linear number of nodes of the tree-decomposition, computing whether there exists a vertex t-ranking costs polynomial time (assuming testing for each applicable value of t (see Lemma 18) for the existence of vertex t-rankings of G, we obtain the following result: Theorem 22 For any fixed k, there exists a polynomial time algorithm, that determines the vertex ranking number of graphs G with treewidth at most k, and finds an optimal vertex ranking of G. 6 The equality - In this section we consider questions related to the equality of the chromatic number and the vertex ranking number of graphs. Theorem 23 If - r holds for a graph G, then G also satisfies Proof: Suppose that E) has a vertex t-ranking with -(G). We are going to consider the separator tree T (c) of this t- ranking. Recall that T (c) is a rooted tree and that every internal node of T (c) is assigned to a subset of the vertex set of G which is a separator of the corresponding subgraph of G, namely more than one component arises when all subsets on the path from the node to the root are deleted from the graph. Furthermore, all vertices assigned to the nodes of a path from a leaf to the root of T (c) have pairwise different colors. The goal of the following recoloring procedure is to show that either !(G) or we can recolor G to obtain a proper coloring with a smaller number of colors. However, the latter contradicts the choice of the -(G)-ranking c. We label the nodes of the tree T (c) according to the following marking rules: 1. Mark a node s of T (c) if the union U(s) of all vertex sets assigned to all nodes on the path from s to the root is not a clique in G. 2. Also, mark a leaf l of T (c) if the union U(l) of all vertex sets assigned to all nodes on the path from l to the root is a clique in G, but jU(l)j ! t. Case 1: There is an unmarked leaf l. We have and U(l) is a clique. Hence, Case 2: There is no unmarked leaf. We will show that this would enable us to recolor G saving one color, contradicting the choice of c. Since every leaf of T (c) is marked, every path from a leaf to the root consists of marked nodes eventually followed by unmarked nodes. Consequently, there is a collection of marked branches of T (c), i.e., subtrees of T (c) induced by one node and all its descendants for which all nodes are marked and the father of the highest node of each branch is unmarked or the highest node is the root of T (c) itself. If the root of T (c) is marked then we have exactly one marked branch, namely T (c) itself. Then, by definition, the separator S assigned to the root is not a clique. However, none of its colors is used by the ranking for vertices in Simply, any coloring of the separator S with fewer than j S j colors will produce a coloring of G with fewer than -(G) colors; contradiction. If the root is unmarked, then we have to work with a collection of b marked branches, b ? 1. Notice that all color-1 vertices of G are assigned to leaves of T (c) and that any leaf of T (c) belongs to some marked branch B. We are going to recolor the graph G by recoloring the marked branches one by one such that the new coloring of G does not use color 1. Let us consider a marked branch B. Let h be its highest node in T (c), and S(h) the set assigned to h. Since h is marked but the root is unmarked, there must exist a vertex x of S(h) and a vertex y belonging to U(h) which are nonadjacent. Then c(x) 6= c(y) since all vertices of U(h) have pairwise different colors. Assume a leaf of T (c). Hence, x and y, respectively, is the only color-1 vertex of G assigned to a node of B. We simply recolor x and y with max(c(x); c(y)). Finally consider the case c(x) 6= 1 and c(y) 6= 1. All color-1 vertices in the subgraph of G corresponding to B are recolored with c(x) and x is recolored with c(y). By the construction of T (c), this does not influence other parts of the graph, since they are separated by vertex sets with higher colors. Having done this operation in every marked branch, eventually we get a new color assignment of G which is still a proper coloring (though usually not a ranking). Since all leaves of T (c) are marked, and no internal node of T (c) contains color-1 vertices, color 1 is eliminated from G, contradicting the assumption - r Consequently, Case 2 cannot occur, implying This completes the proof. 2 does not imply that G is a perfect graph. (Trivial counterexamples are of the form is an arbitrary imperfect graph.) On the other hand, if we require the equality on all induced subgraphs, then we remain with a relatively small class of graphs that is also called 'trivially perfect' in the literature (cf. [11]). Theorem 24 A graph E) satisfies - r if and only if neither P 4 nor C 4 is an induced subgraph of G. Proof: The condition is necessary since - r (P 4 2. Now let G be a P 4 -free and C 4 -free graph. The graphs with no induced are precisely those in which every connected induced subgraph H contains a dominating vertex w, i.e., w is adjacent to all vertices of H [27]. Hence, the following efficient algorithm produces an optimal ranking in such graphs: If then we assign the color !(H) to a dominating vertex w. Clearly, -(H[V 0 it is easily seen that - r (H[V 0 thus, induction can be applied. On the other hand, if H is disconnected, then an optimal ranking can be generated in each of its components separately. 2 7 Edge rankings of complete graphs While obviously - r (K n not easy to give a closed formula for the edge ranking number of the complete graph. The most convenient way to determine r (K n ) seems to introduce a function g(n) by the rules In terms of this g(n), the following statement can be proved. Theorem 25 For every positive integer n, r (K n Proof: The assertion is obviously true for 3. For larger values of n we are going to apply induction. Similarly to vertex t-rankings, the following property holds for every edge t-ranking of a graph is the largest color occurring more than once, then the edges with colors form an edge separator of G. Moreover, doing an appropriate relabeling of these colors get a new edge t-ranking of G with the property that there is a color j ? i such that all edges with colors form an edge separator of G which is minimal under inclusion. We have to show that the best way to choose this edge separator R with respect to an edge ranking in a complete graph is by making the two components of G[E n R] as equal-sized as possible. Let us consider a K n , 2 be the numbers of vertices in the components, hence and the corresponding edge separator has size n 1 n 2 . Every edge ranking starting with this separator has at least r (K n1 r (K r (K maxfn1 ;n2 g ) colors, and there is indeed one using exactly that many colors. Defining a 1 := repeating the same argument for n 0 so on, we eventually get a sequence of positive integers a s, such that a i!j-s a j for all s: (1) Notice that at least the last two terms of any such sequence are equal to 1. It is easy to see that the number of colors of any edge ranking represented by a s is equal to 1-i!j-s a i a j , consequently r (K n 1-i!j-s a i a s a i! subject to the condition (1). Since a decreasing sort of the sequence maintains (1) we may assume a 1 - a 2 - a s . Thus, for each value of 1-i!j-s a i a j is attained precisely by the unique sequence satisfying a i-j-s a j c for all In particular, we obtain r (K n r (K dn=2e Applying this recursion, it is not difficult to verify that, indeed, - 0 r (K n ) can be written in the form 1 is the function defined above. 2 Observing that g(2 n we obtain the following interesting result. Corollary 26 Conclusions We studied algorithmic and graph-theoretic properties of rankings of graphs. For many special classes of graphs, the algorithmic complexity of vertex ranking is now known. However the algorithmic complexity of vertex ranking when restricted to chordal graphs or circle graphs is still unknown. Furthermore it is not even known whether the edge ranking problem is NP-complete. We started a graph-theoretic study of vertex ranking and edge ranking as a particular kind of proper (vertex) coloring and proper edge coloring, respec- tively. Much research has to be done in this direction. It is of particular interest which of the well-known problems in the theory of vertex colorings and edge colorings are also worth studying for vertex rankings and edge rankings. --R A linear time algorithm for finding tree-decompositions of small treewidth A tourist guide through treewidth. Approximating treewidth Graph Theory with Applications. Edge ranking of trees. The multifrontal solution of indefinite sparse symmetric linear equations. Computers and Intractability: A Guide to the Theory of NP-completeness Algorithmic Graph Theory and Perfect Graphs. Optimal node ranking of trees. Parallel assembly of modular products-an analysis On an edge ranking problem of trees and graphs. Ordered colourings. Area efficient graph layouts for VLSI. The role of elimination trees in sparse factorization. Concurrent Design of Products and Processes. The complexity of optimal elimination trees. Graph minors. Graph minors. Graph minors. Node ranking and searching on graphs (Abstract). On a graph partition problem with application to VLSI layout. The comparability graph of a tree. Finding optimal edge-rankings of trees An efficient algorithm for edge-ranking trees --TR --CTR Tak Wah Lam , Fung Ling Yue, Optimal edge ranking of trees in linear time, Proceedings of the ninth annual ACM-SIAM symposium on Discrete algorithms, p.436-445, January 25-27, 1998, San Francisco, California, United States Kazuhisa Makino , Yushi Uno , Toshihide Ibaraki, Minimum edge ranking spanning trees of split graphs, Discrete Applied Mathematics, v.154 n.16, p.2373-2386, 1 November 2006 Shin-ichi Nakayama , Shigeru Masuyama, A polynomial time algorithm for obtaining minimum edge ranking on two-connected outerplanar graphs, Information Processing Letters, v.103 n.6, p.216-221, September, 2007 Keizo Miyata , Shigeru Masuyama , Shin-ichi Nakayama , Liang Zhao, Np-hardness proof and an approximation algorithm for the minimum vertex ranking spanning tree problem, Discrete Applied Mathematics, v.154 n.16, p.2402-2410, 1 November 2006 Md. Abul Kashem , M. Ziaur Rahman, An optimal parallel algorithm for c-vertex-ranking of trees, Information Processing Letters, v.92 n.4, p.179-184, November 2004 Dariusz Dereniowski , Adam Nadolski, Vertex rankings of chordal graphs and weighted trees, Information Processing Letters, v.98 n.3, p.96-100, 16 May 2006 Chung-Hsien Hsu , Sheng-Lung Peng , Chong-Hui Shi, Constructing a minimum height elimination tree of a tree in linear time, Information Sciences: an International Journal, v.177 n.12, p.2473-2479, June, 2007

edge ranking;treewidth;vertex ranking;graph algorithms;ranking of graphs;graph coloring

288976

Directions of Motion Fields are Hardly Ever Ambiguous.

If instead of the full motion field, we consider only the direction of the motion field due to a rigid motion, what can we say about the three-dimensional motion information contained in it? This paper provides a geometric analysis of this question based solely on the constraint that the depth of the surfaces in view is positive. The motivation behind this analysis is to provide a theoretical foundation for image constraints employing only the sign of flow in various directions and justify their utilization for addressing 3D dynamic vision problems.It is shown that, considering as the imaging surface the whole sphere, independently of the scene in view, two different rigid motions cannot give rise to the same directional motion field. If we restrict the image to half of a sphere (or an infinitely large image plane) two different rigid motions with instantaneous translational and rotational velocities(t<math>_1 cannot give rise to the same directional motion field unless the plane through t_1 and t_2 is perpendicular to the plane through _1 and _2 (i.e., (t_1 t_2) (_1 In addition, in order to give practical significance to these uniqueness results for the case of a limited field of view, we also characterize the locations on the image where the motion vectors due to the different motions must have different directions.If (_1 _2) (t_1 t_2) &equals; 0 and certain additional constraints are met, then the two rigid motions could produce motion fields with the same direction. For this to happen the depth of each corresponding surface has to be within a certain range, defined by a second and a third order surface. Similar more restrictive constraints are obtained for the case of multiple motions. Consequently, directions of motion fields are hardly ever ambiguous. A byproduct of the analysis is that full motion fields are never ambiguous with a half sphere as the imaging surface.

addition, in order to give practical significance to these uniqueness results for the case of a limited field of view, we also characterize the locations on the image where the motion vectors due to the different motions must have different directions. additional constraints are met, then the two rigid motions could produce motion fields with the same direction. For this to happen the depth of each corresponding surface has to be within a certain range, defined by a second and a third order surface. Finally, as a byproduct of the analysis it is shown that if we also consider the constraint of positive depth the full motion field on a half sphere uniquely constrains 3D motion independently of the scene in view. The support of the Advanced Research Projects Agency (ARPA Order No. 8459) and the U.S. Army Topographic Engineering Center under Contract DACA76-92-C-0009, the Office of Naval Research under Contract N00014-93-1-0257, National Science Foundation under Grant IRI-90-57934, , and the Austrian "Fonds zur F-orderung der wissenschaftlichen Forschung", project No. S 7003, is gratefully acknowledged. 1 Introduction and Motivation The basis of the majority of visual motion studies has been the motion field, i.e., the projection of the velocities of 3D scene points on the image. Classical results on the uniqueness of motion fields [6, 9, 10] as well as displacement fields [8, 12, 14] have formed the foundation of most research on rigid motion analysis that addressed the 3D motion problem by first approximating the motion field through the optical flow and then interpreting the optical flow to obtain 3D motion and structure [2, 7, 13, 15]. The difficulties involved in the estimation of optical flow have recently given rise to a small number of studies considering as input to the visual motion interpretation process some partial optical flow information. In particular the projection of the optical flow on the gradient direction, the so-called normal flow [5, 11], and the projections of the flow on different directions [1, 3] have been utilized. In [3] constraints on the sign of the projection of the flow on various directions were presented. These constraints on the sign of the flow were derived using only the rigid motion model, with the only constraint on the scene being that the depth in view has to be positive at every point-the so-called "depth-positivity" constraint. In the sequel we are led naturally to the question of what these constraints, or more generally any constraint on the sign of the flow, can possibly tell us about three-dimensional motion and the structure of the scene in view. Thus we would like to investigate the amount of information in the sign of the projection of the flow. Since knowing the sign of the projection of a motion vector in all directions is equivalent to knowing the direction of the motion vector, our question amounts to studying the relationship between the directions of 2D motion vectors and 3D rigid motion. We next state the well-known equation for rigid motion for the case of a spherical imaging surface. We describe the constraints and discuss the information exploited when using the full flow as opposed to the information employed when using only the direction of flow. As will be shown, whereas full flow allows for derivation of the direction of translation and the complete rotation, from the orientation of the flow only the direction of translation and the direction of rotation can be obtained. The 2D motion field on the imaging surface is the projection of the 3D motion field of the scene points moving relative to that surface. Suppose the observer is moving rigidly with instantaneous translation Figure 1); then each scene point measured with respect to a coordinate system OXY Z fixed to the camera moves relative to the camera with velocity - R, where If the center of projection is at the origin and the image is formed on a sphere with radius 1, the relationship between the image point r and the scene point R under perspective projection is R with jRj being the norm of the vector R. R f O Figure 1: Image formation on a spherical retina under perspective projection. If we now differentiate r with respect to time and substitute for - R, we obtain the following equation for - r: (r \Theta (t \Theta The first term v tr (r) corresponds to the translational component which depends on the depth the distance of R to the center of projection. The direction of v tr (r) is along great circles (longitudes) pointing away from the Focus of Expansion (t) and towards the Focus of Contraction (\Gammat). The second term v rot (r) corresponds to the rotational component which is independent of depth. Its direction is along latitudes around the axis of rotation (coun- around ! and clockwise around \Gamma!). See Figure 2a, b and c for translational, rotational, and general motion fields on the sphere. (a) (b) (c) Figure 2: Example of (a) a rotational, (b) a translational, (c) a general motion field on a sphere. As can be seen, without additional constraints there is an ambiguity in the computation of shape and translation. It is not possible to disentangle the effects of t and jRj, and thus we can only derive the direction of translation. If all we have is the direction of the flow we can project - r on any unit vector n i on the image and obtain an inequality constraint: r From this inequality we certainly cannot recover the magnitude of translation, since the optical flow already does not allow us to compute it. In addition we are also restricted in the computation of the rotational parameters. If we multiply ! by a positive constant, leave t fixed, but multiply 1 by the same positive constant, the sign of the flow is not affected. Thus from the direction of the flow we can at most compute the axis of rotation and, as discussed before, the axis of translation. Hereafter, for the sake of brevity, we will refer to the motion field also as the flow field or simply flow, and to the direction of the motion field as the directional flow field or simply directional flow. Between the Orientation of the Flow and the Depth-positivity Constraint If we have the flow - r, we know the value of the projection of - r on any direction and we set all the possible information by choosing two directions and (usually orthogonal). Thus we have r We can solve equation (1) for the depth, r Knowing the value in both directions n 1 and we know that the inverse depth has to be the same, and also has to be positive; thus r r If on the other hand we do not use the value of the flow but only its direction and thus the sign of the projection of the flow on n i , then the only constraint that can be utilized is the inequality, which comes from the fact that the depth is positive. Using only the orientation of the flow we obtain for every direction r This inequality provides inequality constraints on the rotational and translational compo- nents, which are independent of the scene: If we consider the sign of the translational component ((t and the sign of the rotational component (! \Theta r) \Delta n i and assume that each of them is either positive or negative, there are 2 \Theta combinations of signs. But once we know the sign of the flow - r one of these four combinations is no longer possible. This observation has been used in the development of global constraints for 3D motion estimation. Choosing directions n i in particular ways the signs of ( - r patterns of positive and negative areas on the image [3-5]. These patterns, whose location and form encodes information about 3D motion, were successfully used in the recovery of egomotion. In this paper, by pursuing a theoretical investigation of the amount of information present in directional flow fields, we demonstrate the power of the qualitative image measurements already used empirically, and justify their utilization in global constraints for three-dimensional dynamic vision problems. The organization of this paper is as follows: In Section 3 we develop the preliminaries- constraints that will be used in the uniqueness analysis. Given two rigid motions, we study what the constraints are on the surfaces in view for the two motion fields to have the same direction at every point. From these constraints, we investigate for which points of the image one of the surfaces must have negative depth. The locations where negative depth occurs are described implicitly in the form of constraints on the signs of functions depending on the image coordinates and the two three-dimensional motions. The existence of image points whose associated depth is negative ensures that the two rigid motions cannot produce motion fields with the same direction. In Section 4, which contains the main uniqueness proof, we study conditions under which two rigid flow fields could have the same direction at every point on a half sphere (i.e., conditions under which there do not exist points of negative depth), and we visualize the locations of negative depth on the sphere. Section 5 is devoted to the treatment of special cases. As a byproduct of the analysis, in Section 6 we investigate the ambiguity of rigid motion for full flow assuming that depth has to be positive, and show that any two different motions can be distinguished on a hemispherical image from full flow. Section 7 summarizes the results. Appendix A studies whether more than two rigid motions could produce the same directional flow field, and the rest of the Appendices (B-F) describe and prove a number of geometric properties used in the main part of the paper. 3 Critical Surface Constraints Let us assume that two different rigid motions yield the same direction of flow at every point in the image. Let t 1 and ! 1 be translational and rotational velocities of the first motion, and let t 2 be translational and rotational velocities of the second motion. Since from the direction of flow we can only recover the directions of the translation and rotation axes, we assume all four vectors t 1 to be of unit length. Let Z 1 (r) and Z 2 (r) be the mapping points r on the image into the real numbers, that represent the depths of the surfaces in view corresponding to the two motions. In the future we will refer to Z 1 and Z 2 as the two depth maps. In this section we investigate the constraints that must be satisfied by Z 1 and Z 2 in order for the two flow fields to have the same direction. We assume that the two depths are positive, and allow Z 1 or Z 2 to be infinitely large. Thus we assume 1=Z 1 3.1 Notation We start by defining some notation: (2) c denotes the triple product of vectors a, b and c. These functions have a simple geometric meaning. If ! 1 any r. If ! 1 is zero for points r lying on a geodesic passing through . In this case f ! defines the locus of points r where v rot 1 (r), the rotational component of the first motion, is parallel to v rot 2 (r), the rotational component of the second motion. Similarly f t (r) is either zero everywhere, or it is zero for points lying on a geodesic passing through t 1 and t 2 . In this case f t is the locus of points r where v tr 1 (r), the translational component of the first motion, is parallel to v tr 2 (r), the translational component of the second motion. r. If they are non-zero, then g ij (r) is zero at points lying on a second order curve consisting of two closed curves on the sphere, the so-called zero motion contour of motion defines the locus of points where v rot i (r) is parallel to v tr j (r) (see Appendix B). Throughout the paper the functions and the curves defined by their zero crossings will play very important roles. To simplify the notation we will usually drop r and write only f i and g ij where the index i in f i can take values t and !. There is a simple relationship between f i and g ij . Let \Theta r)) \Theta (! 1 \Gammag 11 r \Theta r)) \Theta (! 2 \Gammag 22 r Since we assume (r \Delta 22 Let \Theta r) \Theta (r \Theta (! 2 \Theta Then From equations (3) and (4) we get 3.2 Conditions for ambiguity Assuming that motion and motion give rise to flow fields with the same direction at any point; then there exists - ? 0 such that (r \Theta (t 1 \Theta \Theta (r \Theta (t 2 \Theta \Theta r) (6) By projecting the vector equation (6) on directions t 2 \Theta r and r \Theta (! 2 \Theta r) we obtain two scalar \Theta r) \Delta (t 2 \Theta \Theta r) \Delta (t 2 \Theta r) (7)Z 1 Since - is positive, from (7) and (8) we get constraints on 1 22 where sgn(\Delta) denotes the sign function. Let us define s 1 \Gammag . At any point, f i and g ij are constant, so equations simple constraints on 1 . We call them the s 1 -constraint and the s 0-constraint respectively. Similarly we can project equation (6) on vectors t 1 \Theta r and r \Theta (! 1 \Theta r) and obtain constraints on 1 We define s 2 . Equations (11) and (12) provide constraints on 1 and we thus call them the s 2 -constraint and the s 0-constraint. Let us now interpret these constraints: 1 is always non-negative; thus, if the two motions corresponding depth maps Z 1 and Z 2 produce flow with the same direction, the depth Z 1 must satisfy either 1 or 1 Thus Z 1 has a relationship to the surfaces: and Equations (13) and (14) provide hybrid definitions of scene surfaces. To express the surfaces in scene coordinates R, we substitute in the above equations Dividing (13) by Z(r) and replacing Z(r) 2 by R 2 in (14) we obtain \Theta t 2 \Theta and (! 2 Thus we see that Z 1 is constrained through (15) by a second order surface and through (16) by a third order surface. At some points it has to be inside the first surface and at some points it has to be outside the first surface. In addition, at some points it has to be inside the (a) (b) Figure 3: Two rigid motions (t 1 constrain the possible depth Z 1 of the first surface by a second and a third order surface. The particular surfaces shown in the coordinate system of the imaging sphere, projected stereographically, correspond to the motion configuration of Figure 7. second surface and at some points it has to be outside the second surface. Figure 3 provides a pictorial description of the two surfaces constraining Z 1 Analogous to the above derivation, from equations (11) and (12) we obtain a further second and third order surface pair, which constrain the depth map Z 2 3.3 Interpretation of surface constraints We next describe the s 1 - and s 0-constraints in detail. For convenience, we express these constraints for 1 then the s 1 -constraint is sgn(g 12 does not depend on Z 1 Thus it is either satisfied by any Z 1 , or it cannot be satisfied by any Z 1 If f t 6= 0, then we get the s 1 -constraint constraint is , if sgn(g 22 , if sgn(g 22 , if sgn(g 22 If g 21 0, then the s 0-constraint does not depend on Z 1 22 Since we assume 1 22 is either 0, or sgn(g 22 ). So the s 0-constraint is At each point we have the additional constraint 1 - 0. If all three constraints can be satisfied simultaneously at a point in the image, then there is an interval (bounded or unbounded) of values of Z 1 satisfying them. If the constraints cannot be satisfied, this means that the two flows at this point cannot have the same direction and we say that we have a contradictory point. In the following table we summarize the three constraints on 1 . According to the inequality relationships from the s 1 -constraint and the s 0-constraint we classify the image points into four categories (type I-IV). The table analyzes the general case, at a point where Type s 1 -constraint s 0-constraint 1 solution interval Solution exists if I If some of f i are zero at a point, we may obtain constraints that do not depend on , or equality constraints. In the table above each image point is assigned to one of four categories (see Figure 4 for an example). Whether, for a given image point, there actually exists a value for 1 satisfying the constraints depends on whether the solution interval at that point is empty or not. Thus we classify all image points on the sphere into three categories, A, B, and C, depending on the kind of solution interval that exists for Z 1 :(A) there exists no solution for there exists a solution, and the interval for Z 1 is (B) bounded, or (C) unbounded. In the latter two cases, we can also check whether the interval has a lower bound greater than 0. Figure 4: Classification of image points by s 1 -constraint and s 0-constraint. The classification of a point into one of the categories (I-IV) depends on the signs of f t , , and g 22 . The existence of a solution interval at a point also depends on the signs of f ! and g 12 at that point and also on the relative values of s 1 and s 0, i.e., on the sign of s 1 are polynomial functions of r. To find out where they change sign, it is enough to find points where they are zero. The sign of s 1 since s 1 (r) and s 0(r) do not have to be continuous. However their discontinuities occur at points where f t can change at those points and at points where s 1 Using (5), we can write 22 So we see that sgn(s 1 change only at points where at least one of f i , g ij is zero. At points where g 22 we have the s 1 -constraint 1 and the s 0-constraint 1 thus at these points the depth Z 1 is uniquely defined. Let us consider the implicit curves f i In the general case, these equations describe two geodesics and four zero motion contours. Each of the curves divides the sphere into areas where the solution interval for Z 1 could be different (areas of class A, B, C). However, not every point on the curves separates different areas. Inside any of the areas, all the points have the same classification (for example an infinite solution interval for with a positive lower bound). Figure 5 shows an example of this classification, although the derivation of how to actually obtain the areas where there does not exist a solution is deferred to the next subsection. A Figure 5: Classification of image points according to the solution interval. (The corresponding motion configuration is displayed in Figure 7.) Up to this point we have been discussing only the constraints for Z 1 . Similarly, from (11) and (12) we have at any point the s 2 -constraint , or 1 ), and the We obtain the same curves dividing the sphere into areas such that inside any of the areas, the type of solution interval for Z 2 is the same. Now we can summarize the results. The curves f i separate the sphere into a number of areas. Each of the areas is either contradictory (i.e., containing only contradictory points), or ambiguous (i.e., containing points where the two motion vectors can have the same direction). Two different rigid motions can produce ambiguous directions of flow if the image contains only points from ambiguous areas. There are also two scene surfaces constraining depth Z 1 and two surfaces constraining depth Z 2 . If the depths do not satisfy the constraints, the two flows are not ambiguous. 3.4 Contradictory points In this section we investigate conditions that must be satisfied when a point is contradictory. Since the type of solution for Z 1 and Z 2 depends on the signs of f i and g ij , we want to describe sign combinations that yield a contradiction. We investigate the general case, i.e., we assume f i 6= 0, and g ij 6= 0 and use the resulting constraints in Section 4. Special cases are treated separately in Section 5. There are two simple conditions yielding contradiction for Z 1 , one for the s 1 -constraint and one for the s 0-constraint. There is no solution for Z 1 and s 1 - 0. This happens under the following condition which is derived from equation (9). Similarly, from (10) we get a contradiction if 1 We get similar conditions for Z 2 . There is no solution for Z 2 and s 2 - 0, or ifZ 2 This happens under conditions C 3 and C 4 and We call these four constraints (C 1 Contradictory Point conditions, or CP-conditions for short. Next we show that a point (where f i 6= 0 and g ij 6= 0) is contradictory if and only if at least one of the four conditions is satisfied. Let us assume that conditions (18) and (19) are not satisfied at some point, but we have a contradiction for Z 1 . Then the point must be of type II or III, since there is always a solution for points of type I, and a point of type IV is contradictory only if (18) or (19) holds. For a point of type II, 1 s 0, but (19) is not satisfied, so we have s 0- 0. A contradiction is possible only if s 1 happens when sgn(g 22 sgn(g 22 and s 1 ), from (17) we obtain Thus in this case condition (20) holds. We obtain the same result for points of type III. Since (18) is not satisfied, we have a contradiction is possible only if s happens when sgn(g 22 and and again condition (20) holds. Thus if there is no solution for Z 1 , at least one of conditions (18), (19) and (20) must hold. Similarly if there is no solution for Z 2 , at least one of conditions (18), (20) and (21) must hold. By examination of all the possibilities, we can show that at any point, either none of the CP-conditions holds (and the point is ambiguous), or exactly two of the conditions hold (and the point is contradictory). Antipodal pairs of points In this section we investigate constraints for a point r and its antipodal point \Gammar to be both ambiguous or to be both contradictory. Again we describe a general case, i.e., assume f i 6= 0 and holds either at r, or at \Gammar. We get similar results for the remaining three CP-conditions. Thus both point r and point \Gammar are ambiguous only if Point r and point \Gammar can also both be contradictory. As shown in Appendix F, this happens when 4 The Geometry of the Depth-Positivity Constraint In the last section we found that if the CP-conditions hold at a point on the imaging surface, then one of the depth values has to be negative and thus the point is contradictory. In this section we investigate these constraints further; in particular we would like to know under what conditions two rigid motions cannot be distinguished if our imaging surface is a half sphere or an image plane, and we are interested in studying and visualizing the locations of areas where the CP-conditions are met. Considering as imaging surface the whole sphere, two different rigid motions cannot produce flow of the same direction everywhere. As shown in Section 3.5, two antipodal points r and \Gammar are ambiguous only if (22) holds. Thus for any point on curve the sign of g ij is positive on one side of the curve and negative on the other, there must exist a neighborhood either around r or around \Gammar where there is a contradiction. We are now ready, using the machinery already developed, to study uniqueness properties. As in the previous section, we assume that vectors t 1 , are of unit length. 4.1 Half sphere image: The general case Let us assume that the image is a half of the sphere. Let us also assume that \Theta t 2 We show that under this condition the two rigid motions cannot produce motion fields with the same direction everywhere in the image. Let us consider the projections of ! 1 on a geodesic n connecting t 1 and t 2 Projection onto the geodesic is well defined for all points r such that r \Theta (t 1 \Theta t 2 we assume (24), the projections of both are well defined. The proof is given in parts A and B. A: Let us first assume that one of does not lie on geodesic n. Without loss of generality, let it be ! 1 Figure Possible sign combinations of g 11 and g 12 in the neighborhood of r 1 The projection of ! 1 onto n is \Theta t 2 \Theta t 2 where the sign is chosen so that r 1 is in the image. Then \Theta t 2 \Theta t 2 \Theta t 2 and g 11 the s 0-constraint is sgn( 1 Clearly, this constraint cannot be satisfied, so r 1 is a contradictory point. We can also show that at least one of the areas around point r 1 is contradictory. Point r 1 lies on zero motion contours g 11 If the two contours cross at this point Appendix C shows that g 11 cannot be tangent), we obtain four areas in the neighborhood of r 1 , and all four possible sign combinations of g 11 and g 12 . If we look at points close enough to r 1 (so that f ! does not change sign), then condition (21) is satisfied in one of the areas, and that area is contradictory. For an illustration see Figure 6. B: Now we need to consider the situation where both lie on geodesic n, i.e., us consider point . We know f t (! 1 is parallel to (t i \Theta ! 1 ) is zero. However, from (24) we have \Theta t 2 or g 22 is non-zero at If g 21 cannot be satisfied and ! 1 is a contradictory point. Again, it is not a singular point. The line tangent to g 11 at has direction ! 1 (and 6= 0, since g 21 is perpendicular to n at this point. Since f t is identical to n, curves with all possible sign combinations. Thus in one of the areas, condition (20) holds, and we obtain a contradictory area. If g 22 cannot be satisfied at . Again, at least one area is contradictory, since contour g 12 is perpendicular to n at this point. This concludes the proof that if (24) is satisfied there exist contradictory areas on the half sphere. Section 4.2 discusses the case when (24) is not satisfied. The rest of this section describes properties of the contradictory areas in order to provide a geometric intuition. Just as we projected ! 1 on geodesic n connecting t 1 and t 2 to obtain r 1 , we project ! 2 on n to obtain r 2 , and we project t 1 and t 2 on geodesic l, connecting , to obtain and r 4 (see Figure 7). Point r 2 is at the intersection of f 22 is at the intersection of f is at the intersection of f 22 By the same argumentation as before, at each of the points we can choose two of the contours f passing through the point and we obtain four areas of different sign combinations in the corresponding terms f i and g ij around the point; it can be shown that one of these areas is contradictory because one of the CP-conditions is met. The CP-conditions are constraints on the signs of the terms f i and g ij . Thus the boundaries of the contradictory areas are formed by the curves f As we have shown the contradictory area and its boundaries must contain the points r 1 , and r 4 For some motion configurations the boundaries also might contain t 1 . It can, however, be verified that no neighborhood around t 1 needs to be con- tradictory. It can also be verified, by examining all the possibilities for the signs of terms f i ij in the CP-conditions, that points t 1 , lie inside a contradictory area, since at least one of their neighboring areas is ambiguous. Figures 8 and 9 show the contradictory areas for both halves of the sphere for two different motion configurations. r r r l1222 x Figure 7: Separation of the sphere through curves f Each of t 1 , and r 4 lies at the intersection of three curves. (a) (b) Figure 8: Contradictory areas for both halves of the sphere for the two motions shown in Figure 7. Finally, let us consider the boundaries of the contradictory areas. As defined in Section 3, we allow the depths of the surfaces in view to take any value greater than zero (including infinity). Thus at any point r the motion vector - r could be in the direction of v rot (r), but not in the direction of v tr (r). This allows us to describe the depth values at possible boundaries of a contradictory area: At points on curve f and Z 2 can be infinite, thus boundary points on this curve are not elements of the contradictory area. Boundary points on all other curves (f are contradictory, since one of the depths Z 1 and Z 2 (a) (b) (c) Figure 9: (a) Motion configuration. (b) and (c) Contradictory areas for both halves of the sphere. would have to be zero. 4.2 Half sphere image: The case when (t 1 \Theta t 2 is perpendicular to (! 1 In this section it is shown that there could exist (t 1 \Theta t 2 to (! 1 ), such that there exist no contradictory areas in one hemisphere. First we investigate possible positions of points t 1 on the hemisphere, bounded by equator q. Then we describe additional conditions on the orientation of vectors respect to the hemisphere. As shown in Section 3.5, two antipodal points r and \Gammar can be ambiguous only if (22) holds. Thus if the border of the area defined by (22) intersects q, there will be a contradiction in the image. If curve intersects q at point p, at least one of the areas around p does not satisfy condition (22). Unless all are on the boundary of the hemisphere (and then the motions are not ambiguous), there is a contradictory area in the image (either around or around \Gammap). be the normal to the plane of q. By intersecting the zero motion contour with the border q of the hemisphere (see Appendix D), we find that real solutions for the intersection point are obtained only if A half sphere contains for each of the translation vectors t i and the rotation vectors exactly one of the vectors +t i or \Gammat i and +! i or \Gamma! i . Let us refer to the vectors in the considered hemisphere as ~ t i and ~ From equation (27), taking into account that we see that l ? 0, either if for any ~ (i.e., ~ an angle greater than 90 ffi ), or ( ~ are such that which means that ~ must be close to the border. is perpendicular to f the projections of ! 1 on f and the projections of t 1 and t 2 on . Point r 1 lies at the intersection of all six curves f Any three curves f intersect only in r 1 and one of the points ~ t 1 , or ~ . Furthermore, since all the zero motion contours have to be closed curves on the hemisphere, we conclude that if there exists a contradictory area, it also has to be in a neighborhood of r 1 . It thus suffices to consider all possible sign combinations of terms . It can be verified that, for a hemisphere to contain only ambiguous areas, the two translations have to have the same sign, that is sgn(t 1 Also the two rotations have to have the same sign, i.e., sgn(! 1 Furthermore, the relative positions of t 1 , have to be such that \Theta t 2 Intuitively this means, when rotating in the orientation given by the rotations in order to make f then the order of points t 1 and t 2 on f opposite to the order of points ~ and ~ on (moving along the same direction along \Gamma1, the order of points and \Gammat 2 on f must be the same as the order of points ~ and ~ on In summary, we have shown that two rigid motions could be ambiguous on one hemi- sphere, if (t 1 \Theta t 2 is perpendicular to (! 1 but only if certain sign and certain distance conditions on t 1 , are met. In addition, as shown in Section 3, the two surfaces in view are constrained by a second and a third order surface (as shown in equations (15) and (16)). Figure gives an example of such a configuration. x (a) (b) Figure 10: Both halves of the sphere showing two rigid motions for which there do not exist contradictory areas in one hemisphere. (a) Hemisphere containing only ambiguous areas. (b) Contradictory areas for the other hemisphere. In the next section we discuss the special cases and show that they do not allow for ambiguity. Thus the case of (t 1 \Theta t 2 ) being perpendicular to (! 1 \Theta ! 2 ) is the only case where two motions can produce the same direction of the motion field on a hemisphere. An analysis concerned with ambiguities due to more than two rigid motions is given in Appendix A. 5 Special Cases In previous sections, we assumed that t 1 \Theta t 2 Here we show that if these conditions do not hold, then the two motions are not ambiguous. In Section 3 we assumed all four vectors t 1 to be of unit length. Here the four vectors can also be zero. Thus we have two different motions (i.e. t 1 such that t 1 \Theta t 2 To cover all possible cases we are required to make a minor assumption about the depth and Z 2 for the case where ! 1 . Then we have t 1 , and f From (10) we obtain the constraint 22 ). So at points where g 21 and g 22 have different signs, the only possible solution is 1 Infinite values for both depths in these areas would result in pure rotational flow fields in these areas and thus in an ambiguity. The same kind of ambiguity would occur if we considered the full flow. Therefore it seems reasonable to assume that at least at one point in the areas where g 11 g 22 ! 0, depths Z 1 and Z 2 are not both infinite. Under this assumption there does not exist ambiguity for the case of ! 1 In the following we thus assume Next we provide a lemma that will be of use in the following proofs concerned with special cases as well as in the proof for full flow in Section 6. As in the previous section, let the image be a half sphere with equator q, let n 0 be a unit vector normal to the plane of q. Then equation Z can be satisfied everywhere in the image only if t j \Theta n 0 are non-zero, there are points in the image where \Theta t 2 must be non-zero and geodesic n connecting t 1 and t 2 is well defined. Equation (28) can be satisfied only if the zero motion contour is degenerate, i.e., ! i \Deltat (as in Figure 14b). Then the contour consists of two great circles. One of the circles must be identical to the geodesic n, and the other circle must be identical to q, the border of the image. This is possible only if t j \Theta n 0 Figure Figure 11: If sgn( 1 Z everywhere, the zero motion contour consists of two great circles, one identical to the border of the hemisphere, the other identical to f We now consider two special cases in parts A and B. A: Let us assume that all t i and ! i are non-zero. everywhere. Thus from condition (9) we obtain sgn(g 12 sgn(g 22 four vectors are non-zero, this is possible only if ! 1 So we only need to consider the case ! 1 0, and since we also assume ! 1 , we have . Then at any point in the image, g 2j \Gammag 1j (r). Thus (9) can be satisfied only if sgn( 1 ). According to the lemma, this is possible only if t 2 \Theta n 0 Similarly (11) can be satisfied only if sgn( 1 ). So from the lemma we obtain \Theta n 0 Therefore we have t 1 \Theta t 2 everywhere, and the motions are contradictory. B: If one of the motion parameters is zero, we obtain either a pure translational or a pure rotational flow field. By considering all the possible cases, it can be verified that the two motions are not ambiguous. Here we just consider one of the more difficult cases. Then at any point, g 11 So from (9) we obtain sgn( 1 ), from (11) we have ). From the lemma, this is possible only if t 2 \Theta n 0 \Theta n 0 thus again we obtain t 1 \Theta t 2 and the motions are contradictory. If two of the motions are zero, that is if either t equivalently either two rotational, or one translational and one rotational field, which obviously cannot have the same direction. 6 Ambiguities of the Full Flow Next we investigate the question whether there can be any ambiguities at all if we consider the complete flow. Horn has shown in [6] that two motions can produce ambiguous flow fields only if the observed surfaces are certain hyperboloids of one sheet. We show that if we also consider the depth positivity constraint and if the image is a half of the sphere, then any two different motions can be distinguished. Let the image be a hemisphere bounded by equator q. Let n 0 be a unit vector normal to the plane of q. As in [6], let us assume that a motion (t 1 along with a depth map Z 1 and a motion along with a depth map Z 2 , yield the same flow field. At each point we obtain a vector equation (r \Theta (t 1 \Theta \Theta (r \Theta (t 2 \Theta \Theta r (29) Projecting on directions t 1 \Theta r and t 2 \Theta r, we obtain equations for the two critical \Theta \Theta then according to the lemma in the previous section, these equations can be satisfied everywhere in the image only if ffi! and t 2 \Theta n 0 Thus we we obtain t 1 \Theta t 2 case corresponds to Section 4.5 in [6], that is, to the case when both critical surfaces consist of intersecting planes.) Therefore we are left only with special cases: Ambiguity can occur only if t 1 \Theta t 2 \Theta t 2 0, from constraint (30) we get for any r (ffi! \Theta r) \Delta (t 2 \Theta must be zero. Similarly from constraint (31) we get t Thus we have a pair of rigid motions with different rotations and zero translations. Clearly these two motions are not ambiguous. There is one special case left, At each point we get a vector equation (r \Theta (t 1 \Theta (r \Theta (t 2 \Theta r)) (33) Since we have two different motions and . So the equation can be satisfied only when 1 all points not lying on geodesic n passing through t 1 and t 2 . If we do not allow infinite depth, the motions are not ambiguous. Conclusions In this paper we have analyzed the amount of information inherent in the directions of rigid flow fields. We have shown that in almost all cases there is enough information to determine up to a multiplicative constant both the 3D-rotational and 3D-translational motion from a hemispherical image. Ambiguities can result only if the surfaces in view satisfy certain inequality and equality constraints. Furthermore, for two 3D motions to be compatible the two translation vectors must lie on a geodesic perpendicular to the geodesic through the two rotation vectors. With this analysis we have also shown that visual motion analysis does not necessarily require the intermediate computation of optical flow or exact correspondence. Instead, many dynamic vision problems might be solved with the use of more qualitative flow estimates if appropriate global constraints are found. Appendix Appendix A Ambiguity due to more than two motions In this appendix we investigate whether three or more different rigid motions and their corresponding surfaces could possibly produce the same direction of the motion field on a hemisphere. We present proofs contradicting the ambiguity of almost all combinations of three rigid motions. Let us consider any three different rigid motions ), such that any two of the directional motion fields produced are the same, i.e. j. In the following proofs it will be shown that in general there exist areas in the image where the corresponding depth Z 3 cannot at the same time allow motions (t 1 to produce the same directional flow. Let us consider the intersections of the zero motion contours g ii = 0. In the sequel we consider separately in part A the general case where two of the zero motion contours intersect in at least two points, and in part B the case where any two zero motion contours are tangent to each other. (Appendix E describes the conditions on the motion parameters for two zero motion contours to be tangential.) A: Let us assume that two of the zero motion contours are not tangential; let these be 22 be the intersection point where \Theta t 2 are parallel. Let p 12 be another intersection point where g 11 and g 22 cross. Vectors v rot 1 (p 12 (p 12 are not parallel, and v tr 1 (p 12 (p 12 (p 12 (p 12 positive - 1 or - 2 were negative, point p 12 would be contradictory). Unless g 33 (p 12 we have v tr 3 (p 12 (p 12 Figure 12a shows a possible configuration of the motion vectors at p 12 We next consider the directions of v rot i (r) and v tr i (r) for points r in the neighborhood of p 12 . Let n 0 be a unit vector in the direction v tr 3 (p 11 (p 12 the sign of (v tr i (r) \Theta v rot i changes from inside g ii = 0 to outside g ii = 0 (that is, for example, the angle between v tr 1 (r) and v rot 1 (r) is greater than 180 ffi inside g 11 smaller than 180 ffi outside g 11 vice versa). The sign of (v tr 3 (r) \Theta v rot 3 is the same in a sufficiently small neighborhood around p 12 . Since g 11 22 at there are four neighborhoods around p 12 with all four possible sign combinations of (v tr 1 \Theta v rot 1 and (v tr 2 \Theta v rot 2 . Thus for points r in one of the neighborhoods, in order for (r) to have the same direction as v tr 3 must lie in an interval [a; b], and for v tr 2 (r) to have the same direction as v tr 3 must lie in an interval [c; d], where the intersection of [a; b] and [c; d] is empty. Therefore the three motions cannot give rise to the same direction at r. For an example see Figure 12b. B: We next consider the case where all three zero motion contours are tangent to each other. For the case where not all three are tangential at the same point, using arguments similar to those used before, we prove that there cannot be an ambiguity. For at least two of the zero motion contours, say g 11 22 0, we have that at the intersection point r 12 the two curvature vectors - g 11 of g 11 22 of g 22 sign. Also, the translational and rotational components are such that v tr 1 Figure 13a and b for an illustration). Let n 0 be a unit vector in the direction v tr 3 (r 12 ). At points r in the neighborhood of r 12 we obtain three of the four possible sign combinations for the signs of (v rot 1 (r) \Theta v tr 1 and (v tr 2 (r) \Theta v rot 2 . In both areas outside g 11 outside (r) \Theta v tr 1 (r) \Theta v rot 2 in one of (a) (b) Figure 12: (a) Possible motion configuration at point p 12 . (b) There must exist a neighborhood around with points r, such that for n 0 (p 12 (p 12 )=jv tr 3 (p 12 (p 12 )j, (v rot 1 \Theta v tr 1 \Theta v rot 2 In order for u 3 (v tr 3 (r)), has to be in the sector S 1 and Z 3 has to take values in the interval (0; b]. In order for u 3 (v tr 3 (r)), has to be in the sector 3 has to take values in the interval [c; 1] with b ! c. the areas (v tr 2 (r) \Theta v rot 1 and in the other (v tr 2 (r) \Theta v rot 1 (v tr 3 (r) \Theta v rot 3 doesn't change sign in the neighborhood of r 12 , in one of the two areas the depth Z 3 of the third surface cannot be compatible with both the first and the second motion (see Figure 13c). r 12 tr tr r 12 (a) (b) (c) Figure 13: (a) Intersection of zero motion contours g 11 22 with 22 motion configuration at point r 12 . (c) At point r in one of the areas outside g 11 22 ) the depth of Z 3 cannot be compatible with both (v tr 1 (r)) and (v tr 2 (r)). Thus, in summary we have shown that more than two different rigid motions can hardly ever give rise to the same direction of flow at every point on a hemisphere. The only possible configurations of motions that may be contradictory, provided the surfaces in view satisfy the constraints described in Section 3, are: a: three or more motions such that the corresponding zero motion contours in the same point b: three or more motions, such that all corresponding zero motion contours g ii = 0 are tangential at the same point r 12 , which as described in Appendix E, can occur only if tan tan Appendix motion contours Let us consider the following question: What is the locus of points where the flow due to the given rigid motion can possibly be zero? As in [4] we can show that such points are constrained to lie on a second order curve on the sphere. The flow at point r can be zero only if the rotational and translational components at r are parallel to each other. Let t and ! be translational and rotational velocity of the observer. Then the flow at point r can be zero only if (r \Theta (t \Theta r)) \Theta (! \Theta By simple vector manipulation, from (34) we obtain Since r 6= 0, the flow at point r can be zero only if (! \Theta r) \Delta (t \Theta Equation (36) describes a second order curve on the sphere, which we will call the zero motion contour of the rigid motion (t; !). The zero-motion contour consists of two closed curves on the sphere. As shown in Figure 14, if (! one of the curves contains t and one contains \Gammat 0 and \Gamma! the two curves become great circles, one orthogonal to t, the other orthogonal to !; if (! \Delta one of the two curves and \Gamma! 0 and the other through \Gammat 0 -t -w -t -w (a) (b) (c) Figure 14: The zero motion contour (the locus of points r where - r could be zero) consists of two closed curves on the sphere. Three possible configurations are (a) (! Appendix C Zero motion contours are not tangent To show that zero motion contours g 11 and g 12 are not tangent at r 1 (see Section 4.1, part A), let us compute tangent lines to the contours at r 1 . Let the direction of the line tangent to at r 1 be . The line lies in the plane tangent to the sphere, so Directional derivative of g 11 along must be zero. Let r dg 11 d" \Theta u 1 \Theta r 1 \Theta r 1 \Theta u 1 \Theta u 1 \Theta u 1 must also satisfy Thus we obtain We can compute the tangent direction from (40) and (37) as \Theta ((! 1 Similarly, the direction tangent to g 12 is \Theta ((! 1 lies on geodesic n, we get \Theta u 2 \Theta t 2 \Theta t 2 \Theta (! 1 \Theta r 1 Also \Theta t 2 and \Theta t 2 \Theta (! 1 \Theta r 1 \Theta t 2 \Theta t 2 \Theta u 2 is not zero, and the two zero motion contours cross at point r 1 Appendix D Zero motion contour crossing the border of the image Let the half sphere image be bounded by equator q, let n 0 be a unit vector normal to the plane of q. We would like to know whether the zero motion contour of motion (t; !) intersects equator q. Let us choose a Cartesian coordinate system such that n 0 Points on equator q can be written as [cos OE; sin OE; 0]. Thus the zero motion contour (! \Theta r) \Delta (t \Theta r) = 0 intersects q if equation has a solution. Writing tan OE, we obtain a quadratic equation This equation has a real solution if After some manipulation, we obtain Appendix Intersections of zero motion contours ) be two ambiguous motions. Let us investigate possible intersection points of the zero motion contours of the two motions. Since ambiguity is possible only when \Theta t 2 \Theta t 2 we can choose a Cartesian coordinate system such that \Theta t 2 \Theta t 2 Y In this coordinate system, we can write t 1 Clearly, both zero motion contours pass through point [0; 0; 1]. We would like to know whether this is the only intersection point. were zero, we would have t 1 thus the zero motion contour g 11 would be degenerate. This is not possible if the motions are ambiguous. Thus W 1 6= 0, and similarly are non-zero. Since the zero motion contour does not depend on the size (and direction) of vectors t and !, we can re-scale vectors t multiplying by - i 6= 0 such that Let us consider point z] such that (! \Theta r) \Delta (t \Theta r) = 0. If z 6= 0, point the equation. Thus it is enough to consider two possible sets of points: points of the form points lying in the plane tangent to the sphere at [0; 0; 1]), and points (these points correspond to points at infinity on the tangent plane). A: To obtain the possible intersection points we express the zero motion contours as \Theta r) \Delta \Theta r) \Delta We can compute y from the difference of the two equations as Substituting (54) into (52), we obtain a polynomial equation of degree 4 in x. One solution is (both zero motion contours pass through point [0; 0; 1]). The remaining equation of degree 3 has at least one real solution. another solution Otherwise the two contours intersect in two different points. Since ! 1 6= 0, we know - the two zero motion contours are tangent only if - . If this is the case, we obtain an equation of degree 2 in x. Its discriminant is Thus if j - 2, the two zero motion contours are tangent at [0; 0; 1], but intersect at two other points. B: Now we compute intersection points we obtain equations xy - so such intersection point exists only if In the previous part we have shown that if - , the two zero motion contours have more than one intersection point. So it is enough to check the tangential case here. If the two contours are tangent at [0; 0; 1], from - j. Since t 1 \Theta t 2 6= 0, this is possible only if - and - Writing sin OE, we obtain equation so again there is an intersection point if j - 2. Therefore the two zero motion contours have only one intersection point if - 2. If we denote the intersection point of g 11 and g 22 as r 12 , this can be written as tan tan tan tan where 6 (\Delta; \Delta) denotes the angle between two vectors. This relationship can also be expressed as Furthermore from (55) we obtain the constraint \Theta t 2 Appendix F Antipodal contradictory points Here for the purpose of providing a description of the areas where two motion fields are ambiguous, the conditions are developed for point r and its antipodal point \Gammar both to be contradictory. Clearly, if one of the CP-conditions holds for r, it cannot be true for \Gammar. So if both r and \Gammar are contradictory, two of the conditions must hold at r and the other two at \Gammar. If (18) and (19) hold at r and (20), (21) at \Gammar, we get at r and at \Gammar. Thus If (18) and (20) hold at r and (19), (21) at \Gammar, we get sgn(g 12 at r and 22 ) at \Gammar, so this case cannot occur. If (18) and (21) hold at r and (19), (20) at \Gammar, we get at r and at \Gammar. So we get Thus point r and point \Gammar are both contradictory if and only if --R Optical flow from 1-D correlation: Application to a simple time-to-crash detector Motion and structure from motion from point and line matches. Passive navigation as a pattern recognition problem. On the geometry of visual correspondence. Qualitative egomotion. Motion fields are hardly ever ambiguous. Relative orientation. A computer algorithm for reconstruction of a scene from two projections. Theory of Reconstruction from Image Motion. Critical surface pairs and triplets. Direct passive navigation. Structure from motion using line correspondences. Dynamic aspects in active vision. Uniqueness and estimation of three-dimensional motion parameters of rigid objects with curved surfaces Robust and fast computation of edge characteristics in image sequences. --TR Direct passive navigation Structure from motion using line correspondences Relative orientation Dynamic aspects in active vision Robust and fast computation of edge characteristics in image sequences Qualitative egomotion Optical flow from 1-D correlation Passive navigation as a pattern recognition problem On the Geometry of Visual Correspondence --CTR Tom Svoboda , Tom Pajdla, Epipolar Geometry for Central Catadioptric Cameras, International Journal of Computer Vision, v.49 n.1, p.23-37, August 2002 Fermller , Yiannis Aloimonos, Ambiguity in Structure from Motion: Sphere versus Plane, International Journal of Computer Vision, v.28 n.2, p.137-154, June 1998 Tom Brodsk , Cornelia Fermller , Yiannis Aloimonos, Structure from Motion: Beyond the Epipolar Constraint, International Journal of Computer Vision, v.37 n.3, p.231-258, June 2000

motion field;egomotion;qualitative vision;optic flow

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A Nested FGMRES Method for Parallel Calculation of Nuclear Reactor Transients.

A semi-iterative method based on a nested application of Flexible Generalized Minimum Residual)FGMRES) was developed to solve the linear systems resulting from the application of the discretized two-phase hydrodynamics equations to nuclear reactor transient problems. The complex three-dimensional reactor problem is decomposed into simpler, more manageable problems which are then recombined sequentially by GMRES algorithms. Mathematically, the method consists of using an inner level GMRES to solve the preconditioner equation for an outer level GMRES. Applications were performed on practical, three-dimensional models of operating Pressurized Water Reactors (PWR). Serial and parallel applications were performed for a reactor model with two different details in the core representation. When appropriately tight convergence was enforced at each GMRES level, the results of the semi-iterative solver were in agreement with existing direct solution methods. For the larger model tested, the serial performance of GMRES was about a factor of 3 better than the direct solver and the parallel speedups were about 4 using 13 processors of the INTEL Paragon. Thus, for the larger problem over an order of magnitude reduction in the execution time was achieved indicating that the use of semi-iterative solvers and parallel computing can considerably reduce the computational load for practical PWR transient calculations.

Introduction . The analysis of nuclear reactor transient behavior has always been one of the most difficult computational problems in nuclear engineering. Because the computational load to calculate detailed three-dimensional solutions of the field equations is prohibitive, variations of the power, flow, and temperature distributions are treated approximately in the reactor calculation resulting in considerable conservatism in reactor operation. Some researchers have estimated that using existing methods, the computational load to calculate three-dimensional distributions would exceed a Teraflop per time step. [16] [6] Over the last several years computer speed and memory have increased dramatically and has motivated a rethinking of the limitations of the existing reactor transient analysis codes. Researchers have begun to adapt three-dimensional hydrodynamics and neutron kinetics codes to advanced computer architectures and have begun to investigate advanced numerical methods which can take full advantage of the potential of high performance computing. The overall goal of the work reported here was to reduce the computational burden for three-dimensional reactor core models, thereby enabling high fidelity reactor system modeling. The specific objective was to investigate Krylov Sub-Space methods for the parallel solution of the linear systems resulting from the reactor hydrodynamics equations. The following section provides a brief description of the hydrodynamic model and reactor problem used in the work here. The nested GMRES method and preconditioner developed in the work here are described in section 3 and serial and parallel applications are presented in sections 4 and 5, respectively. This work was supported by the Electric Power Research Institute y School of Nuclear Eng., Purdue University, W. Lafayette, IN 47907 2. Hydrodynamic Model and Reactor Problem. The nuclear reactor analysis problem involves the solution of the coupled neutron kinetics, heat conduction, and hydrodynamics equations. Of these, two-phase flow hydrodynamics is generally the most computationally demanding and was the focus of the work here. The hydrodynamic method used here is consistent with the reactor systems code RETRAN-03 [7] which is widely used in the nuclear industry for the analysis of reactor transient behavior. The method is based on a semi-implicit solution of the mass and momentum equations for each phase and the energy equation for the fluid mixture. The solution scheme uses finite difference representations for all of the fluid-flow balance equations. All convective quantities and most source terms are linearized, i.e. expanded using a first order Taylor series. As a result, the coupling between the finite difference equations is implicitly included in the system of coupled equations. The method is considered only semi-implicit since linearized equations are used rather than the original nonlinear partial differential equations. A standard Newton-Raphson technique is used to solve the nonlinear equations. 2.1. Hydrodynamic Model. The application of spatial finite differencing to the set of governing partial differential equations is performed using the concept of volumes and connecting junctions. This results in a system of ordinary differential- difference equations which may be expressed as dY dt (1) where Y is a column vector of nodal variables and F is a column vector of functions is the number of dependent/solution variables. The solution vector, Y consists of N j junction mass flow rates (W ) and slip velocities (V SL ), and volume total mass (M ), total energy (U ), and vapor mass (M g ) inventories. (2) The vector F(Y) is linearized and a first order time difference approximation is used, resulting in the following linear system: where I is the identity matrix and J is the matrix Jacobian. \DeltaY and \Deltat are the values of Y at time levels t n+1 and t n respectively. Because of the semi-implicit nature of the formulation, the linear system that arises after linearizing and discretizing is tightly coupled. Also, the high degree of stability of this formulation imposes a less stringent limit on the size of the time step. As indicated in Equation 3, larger time steps further reduce diagonal dominance and results in a more ill-conditioned linear system. Several authors have noted [5] [3] that such ill-conditioned linear systems with no well-defined structure are very difficult to solve efficiently and in many cases special handling is required to obtain an acceptable solution. The present linear solver in RETRAN-03 is a direct method based on Gaussian elimination which has a \Theta(N 3 ) execution time complexity where N is the number of unknowns. For models dominated by one-dimensional flow, direct solutions complemented by some type of matrix reduction technique can be very efficient. However, in the case of a high fidelity model of a reactor with a three dimensional core rep- resentation, direct methods become inefficient. This can be demonstrated using the model of a standard 4-loop Presurized Water Reactor. A A A A A A A AAA AAA AAA AAA AAA AAA AAA A A A A A A AAA AAA AAA AAA AAA AAA A A A A A A A AAA AAA AAA AAA AAA AAA AAA A A A A A A AAA AAA AAA AAA AAA AAA System Model Model see Model below Side View Top View channel Upper Plenum 26 33 28 14 31 17 43 23 48 36 343111 355769 7177737579656 4 by pass Fig. 1. Nodalization for the PWR Model 2.2. Pressurized Water Reactor Problem. The reactor system model used in this work consists of 4-loop pressurized water reactor model with a detailed three dimensional representation of the reactor core. A schematic is shown in Figure 1. Each loop of the system contains a steam generator in which heat is transferred from the pressurized primary loop to a secondary loop containing the steam turbines. The core model consists of several volumes stacked one upon the other along each channel. Assuming a layout of C by C channels and A axial volumes, the number of volumes is AC 2 , the number of cross flow junctions is and the number of vertical junctions is C 2 (A+1). The first problem used in the work here had a core model with 9 channels and 12 axial levels per channel as shown in the Figure. We constructed two versions of the core model, one with cross flow (horizontal flow) between volumes and the other without cross flow. Cross flow is important for reactor transients such as the break of a steam line in one of the loops which results in the horizontal mixing of hot and cold water in the core. The model with cross flow consists of 173 volumes and 339 junctions which results in a linear system of size 1173. The core of the model with cross flow contains 108 volumes and 261 junctions and forms nearly 70% of the system. The model without cross flow contains the same number of volumes but has only 183 junctions and results in a linear system of size 885. The coefficient matrix from the problem with cross flow is shown schematically in Figure 2. The structure shown in the figure corresponds to a sequential ordering of the junction and volume variables which is more convenient for the reduction/elimination method currently used in the code. A reordering that is more suitable for preconditioning will be discussed in the following section. Columns Rows Fig. 2. PWR Model with 3-D Core: Sparsity Pattern of the Coefficient Matrix For an initial test problem a simple core reactivity event was modeled by simulating the insertion and withdrawal of a control rod for 20 seconds. The rod was inserted into the core at a rate of 0.08 dollars of reactivity per second for 5 seconds and then immediately withdrawn at the same rate. The computational performance for this problem on a single processor of the INTEL Paragon is summarized in Table 1 for the models with and without cross flow. The direct linear system solution is performed in the SOLVE module indicated in the Table. Table Computatinal Summary for PWR Example Problem with Direct Linear Solver Module w/o Cross Flow w/ Cross Flow CPU Time Percent CPU Time Percent OTHER 155.89 46.7 196.17 4.9 TOTAL 333.88 100.0 4931.71 100.0 It is immediately apparent from Table 1 that the use of cross flow in the core increases the computational time by an order of magnitude. This is primarily because the problem without cross flow results in a linear system containing predominantly tridiagonal submatrices which lends itself very efficiently to reduction/elimination methods. This large increase in the execution time discourages the modeling of cross flow and, in general, high fidelity reactor simulation. Although a reduction in the operation count required for a direct solution was the primary motivation for the work here, there were other reasons to consider iterative linear solution methods. First, the actual problem being solved in RETRAN-03 is nonlinear and the direct solution of the resultant linearized equations can be a waste of floating point operations since usually an accuracy considerably less than machine precision is adequate. Secondly, unlike direct methods, semi-iterative solution methods can be accelerated with information from the previous time steps, such as an initial guess or a preconditioner. Finally, direct methods do not lend themselves easily to parallel computing on distributed memory, MIMD multicomputers, whereas, many of the new semi-iterative linear solvers can be efficiently parallelized. 3. Semi-Iterative Linear Solvers for 3-D Hydrodynamics. During the last several years, considerable research has been performed on a non-stationary class of techniques, collectively known as Krylov subspace methods. These include the classical Conjugate Gradient (CG) method which has been shown to be very efficient for symmetric positive definite systems of equations. These methods are called Krylov methods because they are based on building a solution vector from the Krylov sub- space: span fr is the residual of the initial solution and A is the coefficient matrix. The coefficients of the solution vector in the case of the CG method are based on the minimization of the energy norm of the error. In gen- eral, the linear systems encountered in hydrodynamics problems are not symmetric positive definite and therefore can not be solved using the CG method. Numerous Krylov methods for solving the non-symmetric problem have been proposed over the years and several were considered in this work to include the Generalized Minimal Residual(GMRES)[10] method, the BiConjugate Gradient (BiCG) method [1], the Conjugate Gradient Squared (CGS) method [12], and the BiConjugate Gradient Stabilized (BiCGS) method [15]. Of particular interest in the work here is the GMRES method which solves a minimization problem at each iteration and therefore guarantees a monotonically decreasing residual. It is well known that the convergence rate of the Krylov methods depends on the spectral properties of the coefficient matrix and that proper preconditioning can considerably improve the rate of convergence. Because preconditioning involves some additional cost, both initially and per iteration, there is a trade-off between the cost of implementing the preconditioner and the gain in convergence speed. Since many of the traditional preconditioners have a large sequential component, there is further trade-off between the serial performance of a preconditioner and its parallel efficiency. Several alternative preconditioners were examined in the work here. 3.1. Application of Krylov Methods to Reactor Problem. Previous work on the application of semi-iterative solvers to reactor core hydrodynamic calculations [13] [8] has focused primarily on the solution of the linearized pressure equation resulting from single phase, one-dimensional flow problems. Some of the Krylov methods were found to perform very well on the resulting tridiagonal system of equations. In particular, Turner achieved excellent convergence with the Conjugate Gradient Squared method and an ILU preconditioner. While this work provided useful insight, the linear systems resulting from the two-phase, three-dimensional flow problems are significantly different and the performance of the various Krylov methods was very different. Figure 3 shows the behavior of the absolute residual (L2 norm) for the application to the PWR problem of various Krylov methods with no preconditioner. The performance of Bi-CGSTAB, BICG, and CGS are very irregular (CGS is not shown because the residual increased by several orders of magnitude in the first few iterations and then continues to increase). Only the GMRES method demonstrated acceptable behavior with a monotonically decreasingly residual which is expected because of its basis in a minimization principle. Linear systems from other time steps were examined and behavior similar to that shown in Figure 3 was observed. Although the floating point operation count per iteration was higher for GMRES, particularly for larger numbers of iterations, it was attractive for applications here because of its inherent robustness. Preconditioning techniques were then examined that could improve the convergence behavior of GMRES. 3.2. Domain Decomposition Preconditioning. Knowledge of the physical characteristics of the system is invaluable in choosing a good preconditioner. It is evident from Figure 1 that the core and the ex-core systems interact only at the upper and lower plenum. This suggests a natural way to decompose the problem domain. If the system were to be reordered such that all the solution variables that belonging to the core are placed contiguously, then the structure of the resultant matrix is given by AC U the blocks L and U represent the interactions between the core and the ex-core vari- ables. Because the core and the ex-core interact only at the upper and lower plenum, these matrices have very few non zero elements. Several options exist for using Equation 4 as a preconditioner. One possible preconditioner is to neglect the L and U blocks entirely, resulting in a block Jacobi preconditioner given by AC GMRES Iterations Residual Fig. 3. Figure Performance of Krylov Solvers on PWR 3-D Core Problem The block AE represents the interactions between variables in the ex-core region and the block AC represents the interactions between variables in the core region. Because the preconditioner is Jacobi, the two blocks may be solved individually which is attractive for parallel computing. As noted earlier, the size of the core block is considerably larger than the ex-core block and in a later section methods are discussed for solving the core problem. The block Jacobi domain decomposition preconditioner was applied to the 3-D PWR example problem first using a direct solution of the preconditioner equation. Table 2 contains the results for the linear systems at the first two time steps of the transient described in section 2.2. The first case corresponds to the problem for which results were shown in Figure 3. In addition to the number of iterations, two other measures of the effectiveness of the preconditioner are shown in the Table. First, the condition of the original and preconditioned matrix is shown in the second and third columns. In the fourth column is shown another measure of effectiveness of preconditioning suggested by Dutto[4] which uses the ratio of the Frobenius norm of the remainder matrix, defined as , to the Frobenius norm of original matrix. The Frobenius norm is defined as diagonal As shown in the Table, both measures indicate the Jacobi preconditioner is very effective. For a tolerance of 10 \Gamma6 (relative residual), convergence is achieved in less Table Results of Application of Domain Decomposition to PWR Problem Preconditioner kAkF Iterations 2:19 Gauss Seidel 8:30 2:19 Modified Gauss Seidel 2:19 In the block Jacobi preconditioner the impact of the submatrices L and U is completely neglected while solving the preconditioner equation. Another possibility is to employ a Gauss-Seidel like technique in which the preconditioner equations are solved sequentially: z C z C and - z C and z E respectively. Two approaches were examined for solving these equations as a preconditioner. In one approach, the first equation is solved neglecting the coupling to the other (i.e. - z C or - Because of symmetry, the sequence of solution does not matter and the preconditioner has the form: AC Results of applying this technique are shown as the Gauss-Seidel preconditioner in Table 2. As indicated there is only minor change in the measures of effectiveness of the preconditioning, however, convergence is achieved in fewer iterations. In the second approach an estimate of the ex-core solution, - z E , is formed using the decoupled ex-core equation: AE - and is then used to solve sequentially the core and ex-core equations given by Equation 7. The preconditioner for this approach can be expressed as: This approach is shown in the Table 2 as the Modified Gauss-Siedel preconditioner and, as indicated, little differences are observed in the effectiveness of the precondi- tioning. Because the Block Jacobi method is naturally parallel, it was used in the work here even though the Gauss-Siedel methods showed slightly better numerical performance. For the results shown above the preconditioner equation was solved directly using Gaussian Elimination. Because the ex-core problem is predominantly one-dimensional flow, a direct solution using reduction/elimination proves to be very efficient. Con- versely, the three-dimensional core model does not lend itself to direct solution and the following section examines the use of a second level GMRES to solve the core problem. 3.3. Preconditioning the Core Problem. The original matrix ordering shown in Figure 2 is not conducive to preconditioning the core problem, AC z several alternate orderings were examined. Because the junctions are physically between volumes, reordering the solution vector such that it bears some resemblance to the physical layout would help in decreasing the profile of the matrix. The goal is to increase the density of the matrix in the regions around the diagonal (i.e reduce bandwidth) and then use that portion of the matrix for preconditioning purposes (e.g Block Jacobi, etc. The structure existing in the core was exploited by defining a supernode that consisted of both volumes and junctions. The physical domain was discretized into several supernodes that introduced homogeneity in the structure of the matrix. The supernode could be considered a fractal for representing the smallest unit of structure present in the system. As shown in Figure 4, the supernode for the core problem consists of a volume and 3 junctions. One of the junctions is the vertical junction upstream to the volume and the other two are crossflow junctions leading out of the volume, each in a different direction. It should be noted that the use of supernodes leads to a problem of extra junctions at some of the exterior channels. These junctions are dummy junctions and are represented in the matrix but do not appear in the solution. Hence the size of the system increases but its condition number remains the same. Volume Vertical Junction Crossflow Junction Crossflow Junction Fig. 4. Structure of a Supernode Several orderings of the supernodes were considered (e.g. Channel-wise, Plane- wise, Cuthill-McKee, etc) and ordering planewise was found to be the best for the purpose of preconditioning. Each plane represents a two dimensional grid of the supernodes, and the planes themselves form a one dimensional structure that are linked to only two neighbors, resulting in a block tridiagonal matrix as shown in Figure 5 and Equation 10. As in the case of the outer level preconditioner, several options exist for using Equation 10 as a preconditioner for the inner level GMRES. 12002006001000Columns Rows Fig. 5. Structure of Coefficient Matrix : Plane Wise Ordering A block diagonal preconditioner which neglects planar coupling was considered first. Table 3 shows the results of solving the second level GMRES during the first outer iteration. Four different block sizes were examined with a convergence criterion of 10 \Gamma6 . The two cases shown, A 1 C and A 2 C are taken from different time steps. Table Results of the Application of Jacobi Preconditioning on the Inner Level Preconditioner block size Condition # kRkF kAkF Iterations Condition # kRkF kAkF Iterations The results for the other outer iterations were slightly different since the source and hence the initial residual of the second level GMRES was different. However, the results from other cases were consistent with the general trend shown in Table 3. As expected, larger block sizes reduce the number of iterations. However, the cost of solving each block directly would increase as N 3 which offsets the reduction in iterations. Also, smaller blocks have the advantage of scalability for parallel computing. A domain decomposition scheme incorporating the interactions between the diagonal blocks was also examined. An approximate solution, - z C , was computed by solving the block Jacobi system which neglects the coupling between adjacent planes: z 0 The prime notation is used to distinguish the z C and r C vectors here from those which occur in the outer iteration, Equation 7. The inner GMRES preconditioner equation is then solved: where the preconditioner, MD , is: As indicated in Table 4, this improved preconditioning does reduce the number of iterations, but solving the preconditioner becomes more expensive. The number of iterations for a block size of 81 is reduced by half but more than twice the number of floating point operations are required to form MD . Table Results of the Application of Domain Decomposition Preconditioning on the Inner Level kAkF Iterations block size Preconditioners other than Block Jacobi were investigated for the core problem. Primarily because of the ill-conditioned nature of the coefficient matrix, popular schemes such as SSOR and ILU were ineffective. For example, the incomplete LU factorization scheme was tested on a linear system arising from the 3-D Core prob- lem. Banded were constructed and the Frobenius norm of the ILU preconditioner was compared to the norm of the exact inverse: As shown, the Frobenius norm of the approximate preconditioned system, M \Gamma1 A is over three orders of magnitude larger than the Frobenius norm of A \Gamma1 A, which suggests an ILU preconditioner would not be very effective. 3.4. Nested GMRES. The methods described in the previous sections were implemented in a nested GMRES algorithm which consists of using an "inner" level GMRES to solve the preconditioner equation for the "outer" level GMRES. Such a strategy was suggested by Van der Vorst and demonstrated successfully for several model problems [14]. The preconditioner for the inner level GMRES itself could be solved using a third level GMRES, but in the applications here a direct solver proved most efficient. A schematic of the nested GMRES algorithm is shown in Figure 6. A more physical interpretation of the algorithm is to view the overall problem as being decomposed into three simpler, more manageable problems which are then recombined sequentially by GMRES algorithms. At the highest level we take advantage of the naturally loose coupling between the core and the ex-core components and solve separately the linear systems of the core and ex-core regions. These solutions are then recombined using the highest or "outer" level GMRES. At the second or "inner" level, GMRES is used to solve the 3-D core flow problem where focus is on the coupling between the vertical flow channels in the core. And finally at the third or lowest level, GMRES or a direct solver is used to restore coupling between nodes in a plane. CORE PROBLEM SYSTEM PROBLEM (GMRES I) (GMRES II) (GMRES III or Direct) A e z A c z Fig. 6. Nested GMRES Algorithm for RETRAN-03 Linear System Solution 3.5. Flexible General Minimum Residual (FGMRES). The solution of the preconditioning equation in the GMRES method with another GMRES algorithm poses a potential problem due to the finite precision of the inner level solution. In the preconditioned GMRES algorithm the solution is first built in the preconditioned subspace and then transformed into the solution space: where the matrix is the set of orthonormal vectors. In the case of an iterative solution for the preconditioner, the transformation is only approximate, the extent of which is determined by the convergence criterion. In the case of ill-conditioned matrices the approximations could be especially troublesome and a very tight convergence is required to minimize error propagation. The problem of the inexact transformation was alleviated in the work here by using a slight variant of the GMRES algorithm in which an extra set of vectors is stored and used to update the solution. This modification of the GMRES algorithm was suggested by Saad and is called Flexible General Minimum Residual(FGMRES). [9] This algorithm allows for the complete variation of the preconditioner from one iteration to the next by storing the result of preconditioning each of the basis vectors, while they are being used to further the Krylov subspace. Instead of using Equation 15, the final transformation to the solution subspace is then performed using . where the matrix . The FGMRES algorithm is given in Appendix A and was implemented in the nested GMRES method. 4. Serial Applications. 4.1. Static Problem. The nested GMRES algorithm was first applied to the linear systems arising from the first few time steps of the rod withdrawal/insertion transient. Parametrics were performed on the convergence criterion and number of iterations for each of the levels. A maximum iteration limit was set on the inner GMRES since in some cases the rate of convergence was very slow (sometimes termed "critical slowing down"). The variation of the outer (highest) level residual during the iterations is shown in Figure 7 for different maximum number of inner iterations (miter). The results indicate that the rate of decrease in the residual for substantially greater than for miter = 20. However, as shown in Figure 8, the results of subsequent timesteps indicate that the difference in the rate of decrease in the residual between miter = gradually diminishes. The residual achieved on the inner (second) level GMRES for the timestep corresponding to Figure 8 are shown in Figure 9. For each of the maximum iteration limits, the residual increases after the first few iterations. One possible explanation is that the most desirable search directions for the outer level GMRES become harder to resolve, leading to a degradation of the performance of the inner level. However, the diminished quality of the preconditioning from the inner level GMRES does not appear to have a deleterious effect on the convergence of the outer iteration. Based on these convergence results and those from the analysis of other time steps, a strategy was formulated for the implementation of the nested GMRES in the transient analysis code RETRAN-03. The following section discusses the results of applying GMRES to the transient problem. Outer Iteration Number Outer Residual Fig. 7. Reduction of Outer Residual: First Time Step Outer Iteration Number Outer Residual Fig. 8. Reduction of Outer Residual: Subsequent Time Step Outer Iteration Number Inner Residual Fig. 9. Performance of Inner GMRES: Subsequent Time Step 4.2. Transient Problem. In the time-dependent iterative solution of the hydrodynamics equations, error from incomplete convergence at one time step can be propagated into the coefficient matrix of subsequent time-steps. A preliminary assessment of the effect of convergence criteria on the quality of the solution was performed using a "null" transient in which the steady-state condition is continued for several seconds with no disturbance to the system. Several outer level convergence criteria were investigated and acceptable performance was achieved with a tolerance of 1.0E- 09 in the relative residual. At higher tolerances some minor deviation was observed (e.g. less than 1% relative error) in performance parameters such as the core power level. This provided initial guidance in setting tolerance and iteration limits for the transient problems. The PWR rod withdrawal/insertion transient described in section 2.2 was then analyzed using RETRAN-03 and nested FGMRES. The model with cross flow in the core was studied using the iterative solver for a transient time of 20 seconds which required steps. The performance of the iterative algorithm was analyzed by first varying the number of outer iterations using a direct solver for the inner (e.g. problem, and then by varying both the number of outer and inner GMRES iterations. The results are shown in Table 5 as CPU seconds per time step. The maximum relative residual error during any of the outer iterations is shown in the fourth column of the table. For purposes of comparison, the RETRAN-03 solution from Table 1 which uses a direct solution of the linear system is repeated as case A.1 in Table 5. In the first two cases (A.2 and A.3) the inner two GMRES levels (see Figure 6) have been replaced with a direct solver and GMRES is used only for the outer or highest level iteration. The objective here was to isolate the impact of the number of outer iterations on the Table Serial Performance of RETRAN with GMRES: 9 Channel/ 12 Axial Case Iterations Number of CPU secs/Time Step Case Outer Inner max( krk A.1 Direct Direct \Gamma\Gamma 318 14.89 15.65 - A.3 12 Direct 4 A.4 A.6 quality of the solution. It can be seen from Table 5 that the maximum residual error decreases as the number of outer iterations increases. However, in both cases A.2 and A.3, no significant error was observed in the important physical parameters. When the number of outer iterations was reduced to four, minor deviations began to occur in the solution after 15 seconds of the transient. It should be noted that the execution times for cases A.2 and A.3 are generally comparable to the direct solver. The GMRES algorithm was then employed for both the outer and inner iterations, keeping the direct solution for the innermost (third) level. In order to gain some insight on the relation between convergence of the inner and outer GMRES algorithms, the number of outer and inner iterations were varied as shown in Table 5 for cases A.4, A.5, A.6 and A.7. The accuracy of these solutions was examined for important physical parameters in the solution. The normalized core power and the pressurizer pressure (volume 58 in Figure 1) are plotted versus time in Figure 10 for the Direct solution (A.1) and for GMRES solutions (A.4 and A.7). Some minor deviation is observed in the solution with iterations. The execution times are greater than the direct (A.1) and GMRES outer with direct inner (A.3) solutions. However, as will be discussed in the next section, the algorithm with inner level GMRES is more attractive for larger problems and for parallel computing. 5. Parallel Applications. One of the attractive features of preconditioned Krylov methods is their potential for parallel computing. The emphasis in this work was on the use of a distributed memory parallel architecture and applications were performed on the INTEL Paragon. This section describes the mapping of the nested GMRES onto the Paragon and the execution time reductions achievable for the PWR sample problem. The most natural mapping of processors for the PWR model was one processor to the ex-core and one to each of the 12 planes in the core model. The matrices were striped row wise, implying that L and AE of Equation 4 were stored on the processor that is assigned the ex-core. Ideally, AC and U should be partitioned among the 12 processors but the repartitioning of the data was found to be expensive and hence a copy of both AC and U were maintained on each PE and only the operations were distributed among the processors. One of the primary concerns in distributing data and computation for parallel processing is the communication overhead incurred in transferring data between pro- cessors. The time necessary to perform a transfer consists of two parts, the time necessary to initiate a transfer which is referred to as the latency and the time necessary to actually transfer the data which depends on the amount of data and the CORE POWER RETRAN SOLUTION 20 inner 20 outer (psia) RETRAN SOLUTION 20 inner 20 outer Fig. 10. Results of the RETRAN-03 with GMRES for PWR Transients (GMRES Inner): 40 sec machine bandwidth. The following sections discuss the parallelization of the inner and outer levels of the nested algorithm with special emphasis given to the communication issues. 5.1. Parallelization of the Outer Iteration. One of the dominant operations for the outer iteration is the matrix vector product. At each iteration the product is formed of the matrix A in Eq.1 with the residual vector. This operation can be broken into four parts. The first part involves the product AE v E and involves no communication since all the data resides on the same processor. The next two parts involve U v C and L v E . These involve transfer of parts of the vector between processor 0, which is assigned the ex-core, and the processors 1 and 12 which are assigned the bottommost and the topmost planes in the core, respectively. The fourth part of the matrix vector product involves AC v C and requires communication of processors 1 through 12 with at most two processors. In general the matrix vector product requires communication between selected processors and has a specific pattern. The vector inner products, on the other hand, require global communication which means that every processor requires some information from all the other processors. This is because it is necessary to sum the product of each element of the first vector with the corresponding element from the second vector. The elements of the vectors are first multiplied and each processor forms a partial sum with the elements that reside in its domain. Because only the sum of the products is required and not the individual products, the number of transfers required can be considerably reduced by employing the well known method of tree-summing for which the communication costs vary as dlog 2 (N )e (as opposed to N in the case of all to all communication). The Gram-Schimdt orthogonalization process used in GMRES involves several inner products during each iteration. However, the result of any of these inner products is not dependent on any of the others. Therefore, the partial sums of all of the inner products can be performed at one time, further reducing the number of transfers required [2]. The least square solution in GMRES involves a reduced linear system and does not involve sufficient operations to merit parallelization [11]. Communication related to the least squares problem was avoided entirely by performing the least square solution simultaneously on all the processors. Also, because the estimate of the residue in GMRES is a consequence of the least squares solution, the termination criterion could be evaluated without the need for additional communication. 5.2. Parallelization of the Inner Iteration. Several of the operations required to parallelize the inner iteration were similar to the outer iteration. The matrix AC is striped row wise and the row corresponding to each plane is stored in the corresponding processor. Because the planes are linked only to their immediate neighbors, the matrix vector product requires two sets of data transfer. The vector inner products and the least square solution are treated in the same manner as in the outer level. Since the preconditioner for the second level is a block Jacobi, the preconditioner equation, could be implemented without any data transfers, irrespective of whether it is solved using a direct solver or an iterative solver. 5.3. Results: PWR Model with 9 Channel/ 12 Axial Core. The parallelized version of RETRAN-03 was executed on the Paragon with 13 processors. This section presents the result of the rod withdrawal case for the PWR model with the 9 channel/12 axial volume core. The transient was analyzed for 20 seconds and the nested FGMRES solution was performed using a tolerance of 10 \Gamma11 and an iteration limit of 20 on both the inner and outer iterations. Table Serial Performance of the Second Level: 9 Channel/12 Axial Time step Outer Inner Serial Execution time Precond Other Total Table Parallel Performance of the Second Level: 9 Channel/12 Axial Time step Outer Inner Parallel Execution time Precond Comm Other Total Speedup Tables 6 and 7 show the results of serial and parallel implementation, respectively, of the inner or second level GMRES. The fourth column in both Tables headed "Pre- cond" indicates the time required for the preconditioning of the second level (GMRES III or Direct Solve in Figure 6) which was performed in this application using a direct solver by means of LU factorization. The first timestep involves additional initialization costs and consistent comparisons between the parallel and serial versions of RETRAN were not possible because of differences in code structure for initialization. The high speedup (6.22) in the first iteration of the second timestep is due to LU factorization which is performed concurrently in the parallel implementation. It should be noted that the bulk of the time is spent on the third level and since this part of the code is naturally parallelizable, one would expect high efficiencies. However, since the problem size is relatively small, the communicationoverhead in implementing the remainder of GMRES significantly reduces the efficiency. This would not be the case for larger problems as will be seen in the next section. Tables 8 and 9 show the results of serial and parallel implementation of the outer level GMRES. Again the bulk of the time is spent in the preconditioning. Similar execution times were obtained for other timesteps. Speedup of slightly greater than two were obtained for most of the outer iterations. The execution time for the first four time steps is summarized in Table 10. Again, a speedup on the order of two was achieved. Serial Performance of the Outer Level: 9 Channel/12 Axial Timestep Outer Serial Execution time Core Ex-core Other Total Table Parallel Performance of the Outer Level: 9 Channel/12 Axial Timestep Outer Parallel Execution time Core Ex-core Comm Other Total Speedup 5.4. Results: PWR Model with 25 Channel/ 12 Axial Core. In order to examine the scalability of these results, the same PWR model was used but with instead of 9 channels in the core. The linear system is nearly three times the size of that arising form the 9 Channel/ 12 Axial case. As anticipated, and as shown in Table 11 the time required to solve this linear system using the direct solver was an order of magnitude larger than that for the 9 Channel/ 12 Axial case. (Compare with case A.1 in Table 5.) The problem was first executed using RETRAN with nested FGMRES solver on a single processor on the Paragon. Because the problem was too large to be executed on an ordinary node which has 32MB of RAM, a special "fat" node of the Paragon was used which has additional memory (128 MB) to support larger applications. As shown in Table 11, the results on a single processor were encouraging since more than a factor of 2 improvement was achieved with GMRES compared to the direct solver. The problem was then executed using the parallel version of the nested GMRES solver. The domain decomposition method was exactly the same as the 9 Channel /12 Axial case with 12 processors assigned to the core and 1 to the ex-core. As expected, the parallel efficiency was much better and the speedups were about a factor of 4 with respect to the serial version of GMRES. Table 11 shows the comparison of results for the first two time steps. As indicated in the Table 11, the parallel execution of nested GMRES provides over an order of magnitude reduction in the execution time compared to the serial execution of the direct solver. Furthermore, since the memory requirement per node was less than 32MB, the parallel version could be executed using standard sized nodes of the Paragon. Thus parallelization can be the key to not just execution time reductions but also alleviating the memory constraints for larger problems. 6. Summary and Conclusions. A nested FGMRES method was developed to solve the linear systems resulting from three-dimensional hydrodynamics equations. Applications were performed on practical Pressurized Water Reactor problems with three-dimensional core models. Serial and parallel applications were performed for Overall Performance of the Nested FGMRES: 9 Channel/12 Axial Timestep Number of Execution time Speedup Outers Serial Parallel Table Execution time for the 25 Channel/ 12 Axial case Case Timestep Execution time per Iteration Outer Total per Inner other Timestep Core Ex-core Precond other Direct Solver 1 - 343.310 both a 9 channel and a 25 channel version of the reactor core. When appropriately tight convergence was enforced at each GMRES level, the semi-iterative solver performed satisfactorily for the duration of a typical transient problem. The serial execution time for the 9 channel model was comparable to the direct solver and the parallel speedup on the INTEL Paragon was a factor of 2-3 when using 13 processors. For the 25 channel model, the serial performance of nested GMRES was about a factor of 3 better than the direct solver and the parallel speedups were in the vicinity of 4, again using 13 processors. Thus, for the 25 channel problem over an order of magnitude reduction in the execution time was achieved. The results here indicate that the use of semi-iterative solvers and parallel computing can considerably reduce the computational load for practical PWR transient calculations. Furthermore, the results here indicate that distributed memory parallel computing can help alleviate constraints on the size of the problem that can be exe- cuted. Finally, the methods developed here are scalable and suggest that it is within reach to model a PWR core where all 193 flow channels are explicitly represented. 7. Acknowledgements . The authors appreciate the work of Mr. Jen-Ying Wu in generating the transient results reported in this paper. --R Marching Algorithms for Elliptic Boundary Value Problems. Reducing the Effect of Global Communication in GMRES(m) and CG on Parallel Distributed Memory Computers Direct methods for sparse matrices The Effect of Ordering on Preconditioned GMRES Algorithm Numerical Methods for Engineers and Scientists Supercomputing Applied to Nuclear Reactors An Assessment of Advanced Numerical Methods for Two-Phase Fluid Flow A Flexible InnerOuter Preconditioned GMRES Algorithm GMRES: A Generalized Minimal Residual Algorithm for Solving Nonsymmetric Linear Systems A Comparison of Preconditioned Nonsymmetric Krylov Methods on a Large-Scale MIMD Machine Performance of Conjugate Gradient-like Algorithms in Transient Two-Phase Subchannel Analysis GMRESR: A Family of Nested GMRES Methods Some Computational Challenges of Developing Efficient Parallel Algorithms for Data-Dependent Computation in Thermal-Hydraulics Supercomputer Applications --TR

fluid dynamics;Preconditioned GMRES;parallel computing;nuclear reactor simulation

289829

Stochastic Integration Rules for Infinite Regions.

Stochastic integration rules are derived for infinite integration intervals, generalizing rules developed by Siegel and O'Brien [ SIAM J. Sci. Statist. Comput., 6 (1985), pp. 169--181] for finite intervals. Then random orthogonal transformations of rules for integrals over the surface of the unit m-sphere are used to produce stochastic rules for these integrals. The two types of rules are combined to produce stochastic rules for multidimensional integrals over infinite regions with Normal or Student-t weights. Example results are presented to illustrate the effectiveness of the new rules.

Introduction A common problem in applied science and statistics is to numerically compute integrals in the form with . For statistics applications the function p(') may be an unnormalized unimodal posterior density function and g(') is some function for which an approximate expected value is needed. We are interested in problems where p(') is approximately multivariate normal (' - Nm (-; \Sigma) or multivariate Student-t (' - t m (-; \Sigma)). In these cases, a standardizing transformation in the form can be determined (possibly using numerical optimization), where - is the point where log(p(')) is maximized, \Sigma is the inverse of the negative of the Hessian matrix for log(p(')) at -, and C is the lower triangular Cholesky factor for \Sigma The transformed integrals then take the form w(jjxjj)f(x)dx; or Student-t). If the approximation to p(') is good, then f(x) Accepted for publication in SIAM Journal on Scientific Computing. y Partially supported by NSF grant DMS-9211640. can be accurately approximated by a low degree polynomial in x, and this motivates our construction of stochastic multidimensional polynomial integrating rules for integrals I(f). This type of integration problem has traditionally been handled using Monte-Carlo algorithms (see the book by Davis and Rabinowitz, 1984, and the more recent paper by Evans and Swartz, 1992). A simple Monte-Carlo algorithm for estimating I(f) might use with the points x i randomly chosen with probability density proportional to w(jjxjj). This Monte-Carlo algorithm, which is an importance sampling algorithm for the original problem of estimating E(g), is often effective, but in cases where the resulting f(x) is not approximately constant, the algorithm can have low accuracy and slow convergence. However, an important feature of simple Monte-Carlo algorithms is the availability of practical and robust error estimates. If we let oe E denote the standard error for the sample, then and ff R \Gammaff dt. The new methods that we will describe can be considered a refinement of this Monte-Carlo with importance sampling algorithm. Simple Monte-Carlo with importance sampling results are exact whenever the importance modified integrand is constant, but our methods will be exact whenever the importance modified integrand is a low degree polynomial. Our methods will also provide a robust error estimate from the sample standard error. The new one-dimensional integration rules that we develop are generalizations of the rules derived for the interval [-1,1], with weight by Siegel and O'Brien (1985). Their work extends earlier work by Hammersley and Handscomb (1964), who also considered the construction of stochastic integration rules for finite intervals. Our work is also partly based on work by Haber (1969), who introduced the word "stochastic" for generalized Monte-Carlo rules. Our development of stochastic multidimensional integration rules requires an additional change of variables to a radial-spherical coordinate system. We let Z z t z=1 Z 1w(r)r z t z=1 w(r)jrj The numerical approximations to I(f) that we propose to use will be products of stochastic integration rules for the radial interval (\Gamma1; 1) with weight w(r)jrj m\Gamma1 , and stochastic rules ( of the same polynomial degree ) for the surface of the unit m-sphere. Averages of properly chosen samples of these rules will provide unbiased estimates for I(f), and standard errors for the samples can be used to provide robust error estimates for the I(f) estimates. Our development was partially motivated by the work of De'ak (1990), who used a transformation to a spherical coordinate system combined with random orthogonal transformations to develop a method for computing multivariate normal probabilities, but he did not consider using higher degree rules. The basic radial integration rules that we use are combinations of the symmetric sums h(\Gammaae))=2. A radial rule R(h) takes the form Given points fae i g, the weights fw i g will be determined so that R has polynomial degree 2n + 1. The points fae i g will be randomly chosen so that R is an unbiased estimate for T For fixed points fae i g, the selection of the weights is a standard integration rule construction problem. If we want a degree d rule, it is sufficient that When k is an odd integer, the equation is automatically satisfied because both R and the integration operator are symmetric. Define P (h; r) by Y Now P (h; r) is a even degree Lagrange interpolating polynomial for h, so it follows from standard interpolation theory, that P (h; r)), and the weights fw i g that we need to make R degree are just integrals of the even degree Lagrange basis functions. We have the following theorem: Theorem 1 If the points fae i g are distinct non-negative real numbers and the weights fw i g are defined by Y is a degree 2n We now describe how to choose the points fae i g so that R is an unbiased estimate for T (h). In order to accomplish this, we need to find a joint probability density function p(ae Z 1Z 1::: for any integrable h. We will explicitly show how to do this when and conjecture the general form for p for n - 3. We will let T use the fact thatZ 1r The case randomly with density 2ae . Then we have Z 1C(ae)2ae Z 1ae For degree three rule for T (h) is If we choose ae - 0 randomly with density 2ae m+1 w(ae) Z 1R 3 (ae) 2ae m+1 w(ae) dae Z 1ae Z 1ae Z 1ae m+1 w(ae) Z 1ae A degree five rule for T (h) is We will choose ae - 0 and randomly with joint density where K is determined by the condition 1. We now need to show that EfR 5 (h). There are three terms in R 5 to consider, so we start with the first one, and we find Z 1Z 1(aeffi) 0: For the second term we find Z 1Z 1ae Z 1Z 1ae Z 1ae Z 1ae Now Z 1Z 1ae Z 1Z 1ae so Because R 5 is symmetric in ae and ffi , the last term in R 5 also has expected value T (h)=2, so we have shown that EfR 5 our results in this section with Proposition 1. Proposition 1 If and the points fae i g, the rules R 2n+1 given by (1) with weights given by (2), are chosen with probability density proportional to Y ae m+1 then R is an unbiased degree 2n We have proved this for 2. The form for the probability density for n ? 2 is a conjectured natural generalization of the Siegel and O'Brien Theorem 5.1 (1985). Because of practical problems associated with generating random ae's from this density when n ? 2 we focus on the cases. 3 Stochastic Spherical Integration Rules The spherical surface integrals will be approximated by averages of random rotations of appropriately chosen rules for the spherical surface. Let ~ with z t be an integration rule that approximates an integral of a function s(z) over the surface Um of the unit m-sphere defined by z t z = 1. If Q is an m \Theta m orthogonal matrix then ~ is also an integration rule for s over Um , because Furthermore, if S has polynomial degree d, then so does SQ , because s(Qz) has the same degree as s(z). If Q is chosen uniformly (see Stewart, 1980) and S has polynomial degree d, then SQ is an unbiased random degree d rule for Um . There are many choices that could be used for S. We consider rules given in the book by Stroud (1971, pages 294-296) and the review paper by Mysovskikh (1980, pages 236-237). The rules that we will combine with radial rules have degree 1, 3 or 5, and we now list them. A simple degree 1 rule is is the surface content of Um , and z is any point on Um . A simple degree 3 rule is with the "1" in the j th position. This rule uses 2m values of s(z). A different degree 3 rule (Mysovskikh, 1980) is is the j th vertex of a regular m-simplex with vertices on Um . The degree 3 rule - S 3 is slightly more expensive to use than S 3 , but it leads to an efficient general degree 5 rule (Mysovskikh, 1980) (s(\Gammay The points y j are determined by taking the midpoints of edges of the m-simplex with vertices projecting those midpoints onto the surface of Um . - values of s(z). A degree five rule which extends S 3 (Stroud, 1971, page 294) is where u j is one of the points in the fully symmetric set that is determined by all possible permutations and sign changes of the coordinates of the point (r; 2. 4 Stochastic Spherical-Radial Integration Rules In this section we combine stochastic radial rules with stochastic spherical rules to produce random rules for I(f). There are many ways that this could be done. A natural approach is to form a stochastic product rule SR Q;ae (f) from a spherical surface rule S and a radial rule R. Such a rule takes the form ~ If S and R both have degree d, then SR Q;ae (f) will also have degree d (Stroud, 1971, Theorem 2.3-1). If Q is a uniformly random orthogonal matrix and ae is random chosen with the correct density for R, then SR Q;ae (f) will be an unbiased estimate for I(f). We have the following theorem: Theorem 2 If ae is random with density given by Proposition 1, S has degree 2n+1 and Q is an m \Theta m uniform random orthogonal matrix, then w(jjxjj)f(x)dx whenever f is a degree 2n w(jjxjj)f(x)dx for any integrable f . We give three examples of SR rules. A degree one rule constructed from S 1 and R 1 is ae Here Q is unnecessary, because uniform random vectors z from Um give unbiased rules. A degree three rule constructed from S 3 and R 3 is A degree five rule constructed from - S 5 and R 5 is (w (w with ~ and ~ ae (f), SR 3 Q;ae and - respectively. A sample of one of these rules can be generated, and the sample average used to estimate I(f). The standard error for the sample can be used to provide an error estimate. For comparison purposes with the examples in Section 6, we will use SR 0 (f) to denote the one point rule f(z), with the components of z chosen from Normal(0,1). SR 0 (f) is just the simple Monte-Carlo rule for I(f) with multivariate normal weight. 5 Implementation Details and Algorithms In this section we focus on integrals of the form w(jjxjj)f(x)dx; . For integrals of this type, we have determined explicit formulas for the radial rule weights, along with explicit methods for generating the random radial rule points. We will also discuss the multivariate Student-t weight We first consider the rule SR 1 ae . In the case so 2. Therefore ae The probability density for ae is proportional to ae degrees of freedom. It is a standard statistical procedure to generate a random ae with this density (Monahan, 1987). A standard procedure for generating uniformly random vectors z from Um , consists of first generating x with components x i random from Normal(0,1) and setting z = x=jjxjj. However, this combined procedure for generating random vectors aez must be equivalent to just generating random z from w(jjzjj). Therefore, all we need to do is generate the components z i random from Normal(0,1), and this is a simpler procedure. We propose the following algorithm for random degree one rules: Degree One Spherical-Radial Rule Integration Algorithm 1. Input ffl, m, f and Nmax . 2. 3. Repeat (a) (b) Generate a random x with x i - Normal(0,1). (c) I +D and 4. Output I - I(f), oe V and N . The input ffl is an error tolerance, the input Nmax provides a limit on the time for the algorithm, and the output oe E is the standard error for the integral estimate I. The algorithm computes I and V using a modified version of a stable one-pass algorithm (Chan and Lewis, 1979). The unscaled sample standard error oe E will usually be an error bound with approximately 68% certainty. Users of this algorithm who desire a higher degree of confidence can scale oe E appropriately. For example, a scale factor of 2 increases the certainty level to approximately 95%. The error estimates obtained by scaling oe E with this algorithm (and the other algorithms in this section) should be used with caution for low N values. These error estimates are based on the use of the Central Limit Theorem to infer that the sample averages SR are approximately Normal. A careful implementation of the algorithms in this section could include an Nmin parameter and/or use a larger scale factor for oe E for small N values. For large N , a scaled oe E should provide a robust, statistically sound error estimate, as long as the multivariate normal model adequately represents the tails in the posterior density. Posterior densities with thicker tails are often more efficiently and reliably handled using a multivariate Student-t model. One technique for monitoring this is discussed by Monahan and Genz (1996). If we consider the Student-t weight, we can see that the density for ae is proportional to r a change of variable shows this to be proportional to a Beta( m (see Devroye, 1986, for generating methods), so the random ae's and the uniformly random vectors z from Um , needed for SR 1 can easily be generated. We can also show jU m jT so the formula for is the same as the formula for the multivariate Normal case. By making appropriate changes to line 3(b) and 3(c) of the previous algorithm, a modified algorithm could be produced. Next, we consider the rule SR 3 Q;ae . Integration by parts with and therefore The probability density for ae is proportional to ae m+1 e \Gammaae 2 =2 , a Chi density with m+2 degrees of freedom. We propose the following algorithm for stochastic degree three rules: Degree Three Spherical-Radial Rule Integration Algorithm 1. Input ffl, m, f and Nmax . 2. 3. Repeat (a) (b) Generate a uniformly random orthogonal m \Theta m matrix Q. (c) Generate a random ae - Chi(m (d) For I +D and 4. Output I - I(f), oe V and N . The random orthogonal matrices Q can be generated using a product of appropriately chosen random reflections (see Stewart, 1980). Other methods are discussed by Devroye (1986, p. 607). If we consider the Student-t weight case, then integration by parts shows that T therefore require - ? 2. In this case, SR 3 becomes Further analysis shows that r m+1 proportional to a Beta( m+2 density, so the random ae's for these SR 3 can easily be generated, and by making appropriate changes to lines 3(c) and 3(e) of the previous algorithm, a modified algorithm could be produced. Finally, we consider the rule - Q;ae . For the weight and a little algebra shows In order to develop an algorithm for - Q;ae , we need a set of regular m-simplex unit vertices g. We use the set given in Stroud (1971, page 345, correcting a minor misprint), where v The joint probability density for (ae; ffi ) is proportional to (aeffi) m+1 e \Gamma(ae 2 is not a standard probability density, but there is a transformation to standard densities. Consider the integral and make the change of variables R -=2 Finally, let sin sin Z 1p The function q m+1 is proportional to a standard Beta(m density. The first inner integral has the resulting ae ! ffi and the second has ae ? ffi . Because these cases are both equally likely and - SR 5 is symmetric in ae and ffi , there is no loss of generality in always using ae ! ffi . Therefore, we choose r from a Chi(2m and q from a Beta(m density, and then will be distributed with joint probability density proportional to (aeffi) m+1 e \Gamma(ae 2 We note here that the same changes of variables could also be used to provide a practical method for generating the corresponding ae and ffi for the Siegel and O'Brien (1985) finite interval rules. This question was not addressed in their paper. We propose the following algorithm for stochastic degree five rules: Degree Five Spherical-Radial Rule Integration Algorithm 1. Input ffl, m, f and Nmax . 2. compute the m-simplex vertices fv j g. 3. Repeat (a) (b) Generate a uniformly random orthogonal m \Theta m matrix Q and set f~v j g. (c) Generate a random r - Chi(2m and set (d) For ffl For 4. Output I - I(f), oe V and N . If we consider the Student-t weight, then it can be shown that T (- \Gamma2)(- \Gamma4) T 0 , and we must have 4. In this case, we could also produce a formula for - SR 5 . However, we have not found any easy method for generating the random ae's and ffi 's needed for R 5 , and so we do not consider this further. Anyway, for large -, \Gamma( -+m so the rules that we have already developed for the multivariate Normal weight should be effective. A possibly significant overhead cost for the SR 3 and - SR 5 rules is the generation of the random orthogonal matrices. Using the algorithm given by Stewart (1980), it can be shown that the cost for generating one such matrix Q is approximately 4m 3 =3 floating point operations (flops) plus the cost of generating m 2 =2 Normal(0,1) random numbers. For SR 3 rules the columns of Q are used for the evaluation points for 2m integrand values, so the overhead cost per integrand value is 2m 2 =3 flops plus the cost of generating m=4 Normal(0,1) random numbers. Once an integrand evaluation point is available, we expect the cost for the evaluation of the integrand to be at least O(m), because there are m components for the input variable for the integrand. However, with application problems in statistics, the posterior density is often a complicated expression made up of a combination of standard elementary functions evaluated using the input variable components combined with the problem data (see the second example in the next section). Therefore, if the O(m) integrand evaluation cost is measured in flops, we expect the constant in O(m) to be very large, so that the 2m 2 =3 flops for the generation of the evaluation point for that integrand evaluation should not be significant unless m is very large. For - SR 5 rules the Q overhead cost per evaluation point drops to approximately 2m=3 flops (plus the cost of m=4 Normal variates), and this is not significant compared to the integrand evaluation cost for typical statistics integration problems. We also note here that we need m=2 and respectively, for the rules SR 1 and SR 0 , per integrand evaluation, so the Normal variate overhead is higher for the two lowest degree rules. Overall, except for very simple integrands or large m values, we do not expect the overhead costs for the four rules to be significant compared to the integrand evaluation cost. 6 Examples We begin with a simple example, where The following table of results we obtained using the SR rules: Table Test Results from SR Rules Values I oe E I oe E I oe E I oe E 1000 These results are as expected, with much smaller standard errors for the higher degree rules. For our second example we use a seven dimensional proportional hazards model problem discussed by Dellaportas and Wright (1991, 1992). The posterior density is given by aet . After we first transform ae using x log(ae), we model p(') with a multivariate normal approximation. So we use after computing the mode - and C for log(ae) log(p). We added a scaling constant e 207:19 , to prevent problems with underflow. In the following table we show results from the use of SR rules to approximate I(f 2 ), and expected vales for each of the integration variables. The constant S in the table is a normalizing constant. For each of the respective SR rules we used the computed value of Table 2: Test Results from SR Rules with 120,000 f 2 Values Integrand I oe E I oe E I oe E I oe E For this example, the - SR 3 and - SR 5 rule results have standard errors that are smaller than the SR 0 and SR 1 rule standard errors by factors that are on average about one half. Because the decrease in standard errors is inversely proportional to the square root of the number of samples, approximately four times as much integrand evaluation work would be needed for this problem when using the SR 0 and SR 1 rules to obtain errors comparable to the errors for the - SR 3 and - SR 5 rules. These results are not as good as those for the previous problem, but the higher degree SR rules are still approximately four times more efficient than the lower degree rules. The degree five rule was not better than the degree three rule for this problem. After the standardizing transformation, the problem is apparently close enough to multivariate normal, so that a rule with degree higher than three does not produce better results. We did not find any significant difference in running times needed by the four algorithms for the results in Table 2, and this supports our analysis of the relative importance of overhead costs for the different rules. The two examples in this section are meant to illustrate the use of the algorithms given in this paper. Much more extensive testing is needed in order to carefully compare these algorithms with other methods available for numerical integration problems in applied statistics. For some of the testing work that has been recently done with these methods we refer the interested reader to the paper by Monahan and Genz (1996). Further testing work is still in progress. 7 Concluding Remarks We have shown how to derive low degree stochastic integration rules for radial integrals with normal and Student-t weight functions. We have also shown how these new rules can be combined with stochastic rules for the surface of the sphere, to provide stochastic rules for infinite multivariate regions with multivariate normal and Student-t weight functions. Results from the examples suggest that averages of samples of these rules can provide more accurate integral estimates than simpler Monte-Carlo importance sampling methods. The standard errors from the samples provide robust error estimates for the new rules. --R Computing Standard Deviations: Accuracy Methods of Numerical Integration Positive Imbedded Integration in Bayesian Analysis A Numerical Integration Strategy in Bayesian Analysis Numerical Quadrature and Cubature Some Integration Strategies for Problems in Statistical Inference. Computing Science and Statistics 24 Monte Carlo Methods Stochastic Quadrature Formulas Continuous Univariate Distributions-I An Algorithm for Generating Chi Random Variables A Comparison of Omnibus Methods for Bayesian Computation The Approximation of Multiple Integrals by using Interpolatory Cubature Formulae Unbiased Monte Carlo Integration Methods with Exactness for Low Order Polynomials The Efficient Generation of Random Orthogonal Matrices with An Application to Condition Estimation The Approximate Calculation of Multiple Integrals --TR

multiple integrals;numerical integration;monte carlo;statistical computation

289834

Multigrid Method for Ill-Conditioned Symmetric Toeplitz Systems.

In this paper, we consider solutions of Toeplitz systems where the Toeplitz matrices An are generated by nonnegative functions with zeros. Since the matrices An are ill conditioned, the convergence factor of classical iterative methods, such as the damped Jacobi method, will approach one as the size n of the matrices becomes large. Here we propose to solve the systems by the multigrid method. The cost per iteration for the method is of O(n log n) operations. For a class of Toeplitz matrices which includes weakly diagonally dominant Toeplitz matrices, we show that the convergence factor of the two-grid method is uniformly bounded below one independent of n, and the full multigrid method has convergence factor depending only on the number of levels. Numerical results are given to illustrate the rate of convergence.

Introduction In this paper we discuss the solutions of ill-conditioned symmetric Toeplitz systems A n by the multigrid method. The n-by-n matrices A n are Toeplitz matrices with generating functions f that are nonnegative even functions. More precisely, the matrices A n are constant along their diagonals with their diagonal entries given by the Fourier coefficients of f : \Gamma- Since f are even functions, we have [A n are symmetric. In [10, pp.64-65], it is shown that the eigenvalues - j lie in the range of f('), i.e. min Moreover, we also have lim f(') and lim Department of Mathematics, Chinese University of Hong Kong, Shatin, Hong Kong. Research supported by HKRGC grants no. CUHK 178/93E and CUHK 316/94E. y Institute of Applied Mathematics, Chinese Academy of Science, Beijing, People's Republic of China. z Department of Mathematics, Chinese University of Hong Kong, Shatin, Hong Kong. Consequently, if f(') is nonnegative and vanishes at some points ' 0 2 [\Gamma-], then the condition number is unbounded as n tends to infinity, i.e. A n is ill-conditioned. In fact, if the zeros of f are of order -, then instance [4]. Superfast direct methods for Toeplitz matrices have been developed around 1980. They can solve n-by-n Toeplitz systems in O(n log 2 n) operations, see for instance [1]. However, their stability properties for ill-conditioned Toeplitz matrices are still unclear. Iterative methods based on the preconditioned conjugate gradient method were proposed in 1985, see [11, 13]. With circulant matrices as preconditioners, the methods require O(n log n) operations per iteration. For Toeplitz systems generated by positive functions, these methods have shown to converge superlinearly. How- ever, circulant preconditioners in general cannot handle Toeplitz matrices generated by functions with zeros, see the numerical results in x6. The band-Toeplitz preconditioners proposed in [4, 5] can handle functions with zeros, but are restricted to the cases where the order of the zeros are even numbers. Thus they are not applicable for functions like Classical iterative methods such as the Jacobi or Gauss-Seidel methods are also not applicable when the generating functions have zeros. Since lim n!1 the convergence factor is expected to approach 1 for large n. In [8, 9], Fiorentino and Serra proposed to use multigrid method coupled with Richardson method as smoother for solving Toeplitz systems. Their numerical results show that the multigrid method gives very good convergence rate for Toeplitz systems generated by nonnegative functions. The cost per iteration of the multigrid method is of O(n log n) operations which is the same as the preconditioned conjugate gradient methods. However, in [8, 9], the convergence of the two-grid method (TGM) on first level is only proved for the so-called band - matrices. These are band matrices that can be diagonalized by sine transform matrices. A typical example is the 1-dimensional discrete Laplacian matrix diag[\Gamma1; 2; \Gamma1]. In general, - matrices are not Toeplitz matrices and vice versa. The proof of convergence of the TGM for Toeplitz matrices was not given there. From the computational point of view, the matrix on the coarser grid in TGM is still too expensive to invert. One therefore usually does not use TGM alone but instead applies the idea of TGM recursively on the coarser grid to get down to the coarsest grid. The resulting method is the full multigrid method (MGM). We remark that the convergence of MGM for Toeplitz matrices or for - matrices was not discussed in [8, 9]. In this paper, we consider the use of MGM for solving ill-conditioned Toeplitz systems. Our interpolation operator is constructed according to the position of the first non-zero entry on the first row of the given Toeplitz matrix and is different from the one proposed by Fiorentino and Serra [8, 9]. We show that for a class of ill-conditioned Toeplitz matrices which includes weakly diagonally dominant Toeplitz matrices, the convergence factor of TGM with our interpolation operator is uniformly bounded below 1 independent of n. We also prove that for this class of Toeplitz matrices, the convergence factor of MGM with V -cycles will be level-dependent. One standard way of removing the level-dependence is to use "better" cycles such as the F - or the W -cycles, see [12]. We remark however that our numerical results show that MGM with V -cycles already gives level-independent convergence. Since the cost per iteration is of O(n log n) operations, the total cost of solving the system is therefore of O(n log n) operations. We note that the class of functions that we can handle includes functions with zeros of order 2 or less and also functions such as which cannot be handled by band-Toeplitz preconditioners proposed in [4, 5]. We will also give examples of functions that can be handled by multigrid method with our interpolation operator but not with the interpolation operator proposed in [8]. The paper is organized as follows. In x2, we introduce the two-grid method and the full multigrid method. In x3, we analyze the convergence rate of two-grid method. We first establish in x3.1 the convergence of two-grid method on the first level for the class of weakly diagonally dominant Toeplitz matrices. The interpolation operator for these matrices can easily be identified. Then in x3.2, we consider a larger class of Toeplitz matrices which are not necessarily diagonally dominant. The convergence of full multigrid method is studied in x4 by establishing the convergence of the two- grid method on the coarser levels. In x5, we give the computational cost of our method. Numerical results are given in x6 to illustrate the effectiveness of our method. Finally, concluding remarks are given in x7. Given a Toeplitz system A n we define a sequence of sub-systems on different levels: Here q is the total number of levels with being the finest level. Thus for and are just the size of the matrix A m . We denote the interpolation and restriction operators by I m respectively. We will choose I m+1 The coarse grid operators are defined by the Galerkin algorithm, i.e. Thus, if A m is symmetric and positive definite, so is A m+1 . The smoothing operator is denoted by Typical smoothing operators are the Jacobi, Gauss-Seidel and Richardson iterations, see for instance [3]. Once the above components are fixed, a multigrid cycling procedure can be set up. Here we concentrate on the V -cycle scheme which is given as follows, see [3, p.48]. procedure then u q := begin do i := 1 to - 1 d m+1 := I m+1 e m+1 do Here I nm is the nm -by-n m identity matrix. If we set 2, the resulting multigrid method is the two-grid method (TGM). 3 Convergence of TGM for Toeplitz Matrices In this section, we discuss the convergence of TGM for Toeplitz matrices. We first give an estimate of the convergence factor for Toeplitz matrices that are weakly diagonally dominant. Then we extend the results to a larger class of Toeplitz matrices. Let us begin by introducing the following notations. We say A positive (respectively semi-positive) definite matrix. In particular, A ? 0 means that A is positive definite. The spectral radius of A is denoted by ae(A). For A ? 0, we define the following inner products which are useful in the convergence analysis of multigrid methods, see [12, p.77-78]: Here h\Delta; \Deltai is the Euclidean inner product. Their respective norms are denoted by k 2. Throughout this section, we denote the fine and coarse grid levels of the TGM as the h- and H-levels respectively. For smoothing operator, we consider the damped-Jacobi iteration, which is given by see [3, p.10]. The following theorem shows that kG h k 1 - 1 if ! is properly chosen. Theorem 1 ([12, p.84]) Suppose A ? 0. Let ff be such thatff Then satisfies Inequality (6) is called the smoothing condition. We see from the theorem that the damped-Jacobi method (4) with For a Toeplitz matrix A generated by an even function f , we see from (1) that ae(A) - Moreover, diag(A) is just a constant multiple of the identity matrix. Thus it is easy to find an ff that satisfies (5). In applications where f is not known a priori, we can estimate ae(A) by the Frobenius norm or matrix 1-norm of A. The estimate can be computed in O(n) operations. For TGM, the correction operator is given by with the convergence factor given by k(G h are the numbers of pre- and post-smoothing steps in the MGM algorithm in x2. For simplicity, we will consider only The other cases can be established similarly as we have kG h k 1 - 1. Thus the convergence factor of our TGM is given by kG h T h k 1 . The following theorem gives a general estimate on this quantity. Theorem 2 ([12, p.89]) Let chosen such that G h satisfies the smoothing condition (6), i.e. Suppose that the interpolation operator I h H has full rank and that there exists a fi ? 0 such that min and the convergence factor of the h-H two-level TGM satisfies r Inequality (7) is called the correcting condition. From Theorems 1 and 2, we see that if ff is chosen according to (5) and that the damped-Jacobi method is used as the smoother, then we only have to establish (7) in order to get the convergence results. We start with the following class of matrices. 3.1 Weakly Diagonally Dominant Toeplitz Matrices In the following, we write n-by-n Toeplitz matrix A generated by f as its j-th diagonal as a j , i.e. [A] is the j-th Fourier coefficient of f . Let ID be the class of Toeplitz matrices generated by functions f that are even, nonnegative and satisfy a Given a matrix A 2 ID, let l be the first non-zero index such that a l 6= 0. If a l ! 0, we define the interpolation operator as I h I l2 I l 1 I l I l .2 I l .C C C C C C C C C C A Here I l is the l-by-l identity matrix. If a l ? 0, we define the interpolation operator as I h I l I l I l I l . Theorem 3 Let A 2 ID and l be the first non-zero index where a l 6= 0. Let the interpolation operator be chosen as in (10) or (11) according to the sign of a l . Then there exists a fi ? 0 independent of n such that (7) holds. In particular, the convergence factor of TGM is bounded uniformly below 1 independent of n. Proof: We will prove the theorem for the case a l ! 0. The proof for the case a l ? 0 is similar and is sketched at the end of this proof. We first assume that n according to (10), we have kl. For any e where For ease of indexing, we set e We note that with I h H as defined in (10) and the norm k \Delta k 0 in (3), we have l Thus (7) is proved if we can bound the right hand side above by fihe h ; Ae h i for some fi independent of e h . To do so, we observe that for the right hand side above, we have a 0 l l \Gammae 2il+j e (2i\Gamma1)l+j - a 0 l \Gammae 2il+j e (2i\Gamma1)l+j - a 0 l is the n h -by-n h Toeplitz matrix generated by 1 \Gamma cos l'. Thus min Hence to establish (7), we only have to prove that for some fi independent of e h . To this end, we note that the n h -by-n h matrix A is generated by a j cos j': But by (9), a In particular, by (1) Thus, by (12), we then have 2a l 2a l Hence (7) holds with Next we consider the case where n h is not of the form (2k+1)l. In this case, we let We then embed the vectors e h and e H into longer vectors e ~ h and e ~ H of size n ~ h and n ~ H by zeros. Then since ~ he ~ we see that the conclusion still holds. We remark that the case where a l ? 0 can be proved similarly. We only have to replace the function above by (1 Since in this case, f n h (') - 2a l (1 we then have From this, we get (15) and hence (7) with fi defined as in (16). 3.2 More General Toeplitz Matrices The condition on ID class matrices is too strong. For example, it excludes the matrix However, from (12) and (13), we see that (7) can be proved if we can find a positive number fi independent of n and an integer l such that Since by (14) and (17), we see that (18) holds for any matrices B in ID, we immediately have the following corollary. Corollary 1 Let A be a symmetric positive definite Toeplitz matrix. If there exists a matrix B 2 ID such that A - B. Then (7) holds provided that the interpolation operator for A is chosen to be the same as that for B. More generally, we see by (1) that if the generating function f of A satisfies min for some l, then (18) holds. Thus we have the following theorem. Theorem 4 Let A be generated by an even function f that satisfies (19) for some l. Let the interpolation operator be chosen as in (10) or (11) according to the sign of a l . Then (7) holds. In particular, the convergence factor of TGM is uniformly bounded below 1 independent of the matrix size. It is easy to prove that (19) holds for any even, nonnegative functions with zeros that are of order 2 or less. As an example, consider and T n [1 \Gamma cos '] 2 ID, it follows from Theorem 4 that if the interpolation operator for A is chosen to be the same as that for T n [1 \Gamma cos '], the convergence factor of the resulting TGM will be bounded uniformly below 1. We note that T n [1 \Gamma cos '] is just the 1-dimensional discrete Laplacian: diag[\Gamma1; 2; 1]. Our interpolation operator here is the same as the usual linear interpolation operator used for such matrices, see [3, p.38]. However, we remark that the matrix is a dense matrix. As another example, consider the dense matrix T n [j'j]. Since -j'j - ' 2 on [\Gamma-], we have by Hence T n [j'j] can also be handled by TGM with the same linear interpolation operator used for Convergence Results for Full Multigrid Method In TGM, the matrix A H on the coarse grid is inverted exactly. From the computational point of view, it will be too expensive. Usually, A H is not solved exactly, but is approximated using the TGM idea recursively on each coarser grid until we get to the coarsest grid. There the operator is inverted exactly. The resulting algorithm is the full multigrid method (MGM). In x3, we have proved the convergence of TGM for the first level. To establish convergence of MGM, we need to prove the convergence of TGM on coarser levels. Recall that on the coarser grid, the operator A H is defined by the Galerkin algorithm (2), i.e. h A h I h H . We note that if n will be a block-Toeplitz- Toeplitz-block matrix and the blocks are l-by-l Toeplitz matrices. In particular, if l = 1, then A H is still a Toeplitz matrix. However, if n h is not of the form (2k will be a sum of a block-Toeplitz-Toeplitz-block matrix and a low rank matrix (with rank less than or equal to 2l). We will only consider the case where n j. For then on each level . Hence the main diagonals of the coarse-grid operators A m , will still be constant. Recall that from the proof of Theorem 3 that (18) implies (7). We now prove that if (18) holds on a finer level, it holds on the next coarser level when the same interpolation operator is used. Theorem 5 Let a h 0 and a H 0 be the main diagonal entries of A h and A H respectively. Let the interpolation operator I h H be defined as in (10) or (11). Suppose that A h - a hfi h for some fi h ? 0 independent of n. Then with a Proof: We first note that if we define the (n H I l I l I l I lC A ; (24) then there exists a permutation matrix P such that I h I nH (cf (10) and (11)). Moreover, for the same permutation matrix P , we have I nH \UpsilonK t I nH+l By (2) and (21), we have h A h I h a hfi h I H But by (25) and (26), we have a hfi h I H I nH \UpsilonK t I nH+l !/ I nH By the definition of K in (24), we have a hfi h Combining this with (28), we get a hfi h I H Hence (27) implies (22) with (23). Recall by (5) that we can choose ff h such that Notice that K t K - I n h and therefore h A h I h I h Thus on the coarser level, we can choose ff H as According to (8), (30) and (23), we see that s s ff h a H s Recursively, we can extend this result from the next coarser-level to the q-th level and hence obtain the level-dependent convergence of the MGM: s s We remark that this level-dependent result is the same as that of most MGM, see for instance [12, 2]. One standard way to overcome level-dependent convergence is to use "better" cycles such as the F - or W -cycles, see [12]. We note however that our numerical results in x6 shows that MGM with V -cycles already gives level-independent convergence. We remark that we can prove the level-independent convergence of MGM in a special case. Theorem 6 Let f(') be such that for some integer l and positive constants c 1 and c 2 . Then for any 1 - m - q, r Proof: From (31), we have Recalling the Galerkin algorithm (2) and using (29) recursively, we then have By the right-hand inequality and (18), we see that c 1 a m: and hence by the left hand side of (32) Therefore by the definition of ff in (5), we see that c 2 a m: According to (8), we then conclude that s r As an example, we see that MGM can be applied to T n [' 2 ] with the usual linear interpolation operator and the resulting method will be level-independent. Computational Cost Let us first consider the case where j. Then on each level, n some k. From the MGM algorithm in x2, we see that if we are using the damped-Jacobi method (4), the pre-smoothing and post-smoothing steps become Thus the main cost on each level depends on the matrix-vector multiplication A m y for some vector y. If we are using one pre-smoothing step and one post-smoothing step, then we require two such matrix vector multiplications - one from the post-smoothing and one from the computation of the residual. We do not need the multiplication in the pre-smoothing step since the initial guess u m is the zero vector. On the finest level, A is a Toeplitz matrix. Hence Ay can be computed in two 2n-length FFTs, see for instance [13]. If l = 1, then on each coarser level, A m will still be a Toeplitz matrix. Hence A m y can be computed in two 2nm -length FFTs. When l ? 1, then on the coarser levels, A m will be a block-Toeplitz-Toeplitz-block matrix with l-by-l Toeplitz sub-blocks. Therefore A m y can also be computed in roughly the same amount of time by using 2-dimensional FFTs. Thus the total cost per MGM iteration is about eight 2n-length FFTs. In comparison, the circulant-preconditioned conjugate gradient methods require two 2n-length FFTs and two n-length FFTs per iteration for the multiplication of Ay and C \Gamma1 y respectively. Here C is the circulant preconditioner, see [13]. The band-Toeplitz preconditioned conjugate gradient methods require two 2n-length FFTs and one band-solver where the band-width depends on the order of the zeros, see [4]. Thus the cost per iteration of using MGM is about 8/3 times as that required by the circulant preconditioned conjugate gradient methods and 4 times of that required by the band-Toeplitz preconditioned conjugate gradient methods. Next we consider the case when n is not of the form 1)l. In that case, on the coarser level, A m will no longer be a block-Toeplitz-Toeplitz-block matrix. Instead it will be a sum of such a matrix and a low rank matrix (with rank less than 2l). Thus the cost of multiplying A m y will be increased by O(ln). 6 Numerical Results In this section, we apply the MGM algorithm in x2 to ill-conditioned real symmetric Toeplitz systems A n We choose as solution a random vector u such that 0 - 1. The right hand side vector b is obtained accordingly. As smoother, we use the damped-Jacobi method (4) with chosen as a 0 =max f(') for pre-smoother and post-smoother. We use one pre-smoothing and one post-smoothing on each level. The zero vector is used as the initial guess and the stopping criterion is kr where r j is the residual vector after j iterations. In the following tables, we give the number of iterations required for convergence using our method, see column under M . For comparison, we also give the number of iterations required by the preconditioned conjugate gradient method with no preconditioner (I), the Strang (S) circulant preconditioner, the T. Chan (C) circulant preconditioner and also the band (B) preconditioners, see [6, 7, 4]. The double asterisk * signifies more than 200 iterations are needed. For the first example, we consider functions with single zero at the point The functions we tried are We note that T n ID. However, we have Therefore, according to Corollary 1, we can use the interpolation operator (10) with Table 1: Number of Iterations for Functions with Single Zero. Next we consider two functions with jumps and a single zero at the point and We note that both matrices T n [J j (')], are not in ID. However, since J j (') - for all ', we can still use the interpolation operator defined by 1 \Gamma cos ' for both T n [J j (')]. We remark that for J 2 ('), since the zero is not of even order, band circulant preconditioners cannot be constructed. Table 2: Number of Iterations for Functions with Jumps. Finally, we consider two functions with multiple zeros. They are and ID. But we note that ' 2 (- both matrices can use the interpolation operator in (10) with l = 2. In particular, our interpolation operator will be different from that proposed in [8], which in this case will use the interpolation operator in (10) with l = 1. Their resulting MGM converges very slowly with convergence factor very close to 1 (about 0.98 for both functions when 64). For comparison, we list in Table 3, the number of MGM iterations required by such interpolation operator under column F . Table 3: Number of Iterations for Functions with Multiple Zeros. 7 Concluding Remarks We have shown that MGM can be used to solve a class of ill-conditioned Toeplitz matrices. The resulting convergence rate is linear. The interpolation operator depends on the location of the first non-zero diagonals of the matrices and its sign. Here we have only proved the convergence of multigrid method with damped-Jacobi as smoothing operator. However, our numerical results show that multigrid method with some other smoothing operators, such as the red-black Jacobi, block-Jacobi and Gauss-Seidel methods, will give better convergence rate. As an example, for the function convergence factors of MGM with the point- and block-Jacobi methods as smoothing operator are found to be about 0:71 and 0:32 respectively for 64 - n - 1024. Acknowledgment We will like to thank Prof. Tony Chan and Dr. J. Zou for their valuable comments. --R Superfast Solution of Real Positive Definite Toeplitz Systems Convergence Estimates for Multigrid Algorithms without Regularity Assumptions Toeplitz Preconditioners for Toeplitz Systems with Nonnegative Generating Func- tions Fast Toeplitz Solvers Based on Band-Toeplitz Preconditioner Toeplitz Equations by Conjugate Gradients with Circulant Precondi- tioner An Optimal Circulant Preconditioner for Toeplitz Systems Multigrid Methods for Toeplitz Matrices Multigrid Methods for Symmetric Positive Definite Block Toeplitz Matrices with Nonnegative Generating Functions Linear and Nonlinear Deconvolution Problems in Multigrid Methods A Proposal for Toeplitz Matrix Calculations --TR --CTR S. Serra Capizzano , E. Tyrtyshnikov, How to prove that a preconditioner cannot be superlinear, Mathematics of Computation, v.72 n.243, p.1305-1316, July

toeplitz matrices;multigrid method

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Accelerated Inexact Newton Schemes for Large Systems of Nonlinear Equations.

Classical iteration methods for linear systems, such as Jacobi iteration, can be accelerated considerably by Krylov subspace methods like GMRES @. In this paper, we describe how inexact Newton methods for nonlinear problems can be accelerated in a similar way and how this leads to a general framework that includes many well-known techniques for solving linear and nonlinear systems, as well as new ones. Inexact Newton methods are frequently used in practice to avoid the expensive exact solution of the large linear system arising in the (possibly also inexact) linearization step of Newton's process. Our framework includes acceleration techniques for the "linear steps" as well as for the "nonlinear steps" in Newton's process. The described class of methods, the accelerated inexact Newton (AIN) methods, contains methods like GMRES and GMRESR for linear systems, Arnoldi and JacDav{} for linear eigenproblems, and many variants of Newton's method, like damped Newton, for general nonlinear problems. As numerical experiments suggest, the AIN{} approach may be useful for the construction of efficient schemes for solving nonlinear problems.

Introduction . Our goal in this paper is twofold. A number of iterative solvers for linear systems of equations, such as FOM [23], GMRES [26], GCR [31], Flexible GMRES [25], GMRESR [29] and GCRO [7], are in structure very similar to iterative methods for linear eigenproblems, like shift and invert Arnoldi [1, 24], Davidson [6, 24], and Jacobi-Davidson [28]. We will show that all these algorithms can be viewed as instances of an Accelerated Inexact Newton (AIN) scheme (cf. Alg. 3), when applied to either linear equations or to linear eigenproblems. This observation may help us in the design and analysis of algorithms by "transporting" algorithmic approaches from one application area to another. Moreover, our aim is to identify efficient AIN schemes for nonlinear problems as well, and we will show how we can learn from the algorithms for linear problems. To be more specific, we will be interested in the numerical approximation of the solution u of the nonlinear equation (1) where F is some smooth (nonlinear) map from a domain in R n (or C n ) that contains the solution u, into R n (or C n ), where n is typically large. Some special types of systems of equations will play an important motivating role in this paper. The first type is the linear system of equations (2) Mathematical Institute, Utrecht University, P. O. Box 80.010, NL-3508 Utrecht, The Nether- lands. E-mail: [email protected], [email protected], [email protected]. D. R. FOKKEMA, G. L. G. SLEIJPEN, AND H. A. VAN DER VORST where A is a nonsingular matrix and b, x are vectors of appropriate size; A and b are given, x is unknown. The dimension n of the problem is typically large and A is often sparse. With is equivalent to (1). This type will serve as the main source of inspiration for our ideas. The second type concerns the generalized linear eigenproblem With we have that, for normalized v, with F (u) := Av \Gamma -Bv, equation (3) is equivalent to (1). This type is an example of a mildly nonlinear system and will serve as an illustration for the similarity between various algorithms, seen as instances of AIN (see Section 6). However, the AIN schemes, that we will discuss, will be applicable to more general nonlinear problems, like, for instance, equations that arise from discretizing nonlinear partial differential equations of the form where\Omega is a domain in R 2 or R 3 , and u satisfies suitable boundary conditions. An example of (4) is, for instance where\Omega is some domain in R 2 and @\Omega (see also Section 8). Guided by the known approaches for the linear system (cf. [25, 29, 7]) and the eigenproblem (cf. [28, 27]) we will define accelerated Inexact Newton schemes for more general nonlinear systems. This leads to a combination of Krylov subspace methods for Inexact Newton (cf. [16, 4] and also [8]) with acceleration techniques (as in [2]) and offers us an overwhelming choice of techniques for further improving the efficiency of Newton type methods. As a side-effect this leads to a surprisingly simple framework for the identification of many well-known methods for linear-, eigen-, and nonlinear problems. Our numerical experiments for nonlinear problems, like problem (5), serve as an illustration for the usefulness of our approach. The rest of this paper is organized as follows. In Section 2 we briefly review the ideas behind the Inexact Newton method. In Section 3 we introduce the Accelerated Inexact Newton methods (AIN). We will examine how iterative methods for linear problems are accelerated and we will distinguish between a Galerkin approach and a Minimal Residual approach. These concepts are then extended to the nonlinear case. In Section 4 we make some comments on the implementation of AIN schemes. In Section 5 we show how many well-known iterative methods for linear problems fit in the AIN framework. In Section 6 and Section 7 we consider instances of AIN for the mildly nonlinear generalized eigenproblem and for more general nonlinear problems. In Section 8 we present our numerical results and some concluding remarks are in Section 9. ACCELERATED INEXACT NEWTON SCHEMES 3 1. choose an initial approximation u 0 . 2. Repeat until u k is accurate (b) Compute the residual r (c) Compute an approximation J k for the Jacobian F 0 (u k ). (d) Solve the correction equation (approximately). Compute an (approximate) solution p k of the correction equation Update. Compute the new approximation: Algorithm 1: Inexact Newton 2. Inexact Newton methods. Newton type methods are very popular for solving systems of nonlinear equations as, for instance, represented by (4). If u k is the approximate solution at iteration number k, Newton's method requires for the next approximate solution of (1), the evaluation of the Jacobian J k := F 0 and the solution of the correction equation Unfortunately, it may be very expensive, or even practically impossible to determine the Jacobian and/or to solve the correction equation exactly, especially for larger systems. In such situations one aims for an approximate solution of the correction equa- tion, possibly with an approximation for the Jacobian (see, e.g., [8]). Alg. 1 is an algorithmical representation of the resulting Inexact Newton scheme. If, for instance, Krylov subspace methods are used for the approximate solution of the correction equation (6), then only directional derivatives are required (cf. [8, 4]); there is no need to evaluate the Jacobian explicitly. If v is a given vector then the vector F 0 (u)v can be approximated using the fact that The combination of a Krylov subspace method with directional derivatives combines the steps 2c and 2d in Alg. 1. For an initial guess u 0 sufficiently close to a solution, Newton's method has asymptotically at least quadratic convergence behavior. However, this quadratic convergence is usually lost if one uses inexact variants and often the convergence is not much faster than linear. In the next section we make suggestions how this (linear) speed of convergence may be improved. Note: It is our aim to restore, as much as possible, the asymptotic convergence behavior of exact Newton and we do not address the question of global convergence. 4 D. R. FOKKEMA, G. L. G. SLEIJPEN, AND H. A. VAN DER VORST 1. choose an initial approximation x 0 . 2. Repeat until x k is accurate (c) (d) x Algorithm 2: Jacobi Iteration 3. Accelerating Inexact Newton methods. Newton's method is a one step method, that is, in each step, Newton's method updates the approximate solution with information from the previous step only. However, in the computational process, subspaces have been built gradually, that contain useful information concerning the problem. This information may be exploited to improve the current approximate solution, and this is what we propose to do. More precisely, we will consider alternative update strategies for step 2e of the Inexact Newton algorithm Alg. 1. 3.1. Acceleration in the linear case. The linear system can be written as F (x) := and \GammaA is the Jacobian of F . When the approximate solution p k is computed as preconditioning matrix (approximating A), then the Inexact Newton algorithm, Alg. 1, reduces to a standard Richardson- type iteration process for the splitting R. For instance, the choice leads to Jacobi iteration (see Alg. 2). One may improve the convergence behavior of standard iterations schemes by ffl using more sophisticated preconditioners M , and/or ffl applying acceleration techniques in the update step. Different preconditioners and different acceleration techniques lead to different algo- rithms, some of which are well-known. Examples of iterations schemes that use more sophisticated preconditioners are, for instance, Gauss-Seidel iteration, where with L is the strict lower triangular part of A and is a relaxation parameter. Examples of iterations schemes that use acceleration techniques are algorithms that take their updates to the approximate solution as a linear combination of previous directions p j . Preferable updates ~ are those for which b \Gamma Ax k+1 , where x k , is minimal in some sense: e.g., kb \Gamma Ax k+1 k 2 is minimal, as in GMRES [26] and GCR [31], or b \Gamma Ax k+1 is orthogonal to the p j for j - k, as in FOM or GENCG [23], or b \Gamma Ax k+1 is "quasi-minimal", as in Bi-CG [17], and QMR [11]. Of course the distinction between preconditioning and acceleration is not a clear one. Acceleration techniques with a limited number of steps can be seen as a kind of dynamic preconditioning as opposed to the static preconditioning with fixed M . In this view one is again free to choose an acceleration technique. Examples of such iteration schemes are Flexible GMRES [25], GMRESR [29] and GCRO [7]. All these accelerated iteration schemes for linear problems construct approximations the solution of a smaller or an ACCELERATED INEXACT NEWTON SCHEMES 5 easier projected problem. For example, GMRES computes y k such that equivalently (AV k ) (b \Gamma such that V as the solution of a larger tri-diagonal problem obtained with oblique projections. For stability (and efficiency) reasons one usually constructs another basis for the span of V k with certain orthogonality properties, depending on the selected approach. 3.2. Acceleration in the nonlinear case. We are interested in methods for finding a zero of a general nonlinear mapping F . For the linear case, the methods mentioned above are, apart from the computation of the residual, essentially a mix of two components: (1) the computation of a new search direction (which involves the residual), and (2) the update of the approximation (which involves the current search direction and possibly previous search directions, and the solution y k of a projected problem). The first component may be interpreted as preconditioning, while the second component is the acceleration. Looking more carefully on how y k in the linear case is computed, we can distinguish between two approaches based on two different conditions. With G k (y) := F the FOM and the other "oblique" approaches lead to methods that compute y such that (for appropriate W k the GMRES approach leads to methods that compute y such that kG k (y)k 2 is minimal (a Minimal Residual condition). From these observations for the linear case we now can formulate iteration schemes for the nonlinear case. The Inexact Newton iteration can be accelerated in a similar way as the standard linear iteration. This acceleration can be accomplished by updating the solution by a correction ~ k in the subspace spanned by all correction directions p j (j - k). To be more precise, the update ~ k for the approximate solution is given by ~ y, where of the search space V k spanned by Furthermore, with G k we propose to determine y by ffl a Galerkin condition on G k (y): y is a solution of where W k is some matrix of the same dimensions as V k , ffl or a Minimal Residual (MR) condition on G k (y): y is a solution of min y ffl or a mix of both, a Restricted Minimal Residual (RMR) condition on G k (y): y is a solution of min y Equation (7) generalizes the FOM approach, while equation (8) generalizes the GMRES approach. Solving (7) means that the component of the residual r k+1 in the subspace W k (spanned by W k ) vanishes. For W k one may choose, for instance, W (as in 6 D. R. FOKKEMA, G. L. G. SLEIJPEN, AND H. A. VAN DER VORST FOM), or W is the component of J k p k orthogonal to W (as in GMRES: w linear equations the Minimal Residual and the Galerkin approach coincide for the last choice. As is known from the linear case, a complication of the Galerkin approach is that equation (7) may have no solution, which means that this approach may lead to breakdown of the method. In order to circumvent this shortcoming to some extend we have formulated the Restricted Minimal Residual approach (9). Compared to (7) this formulation is also attractive for another reason: one can apply standard Gauss-Newton [9] schemes for solving general nonlinear least squares problems to it. One might argue that a drawback of a Gauss-Newton scheme is that it may converge slowly (or not at all). However, for least squares problems with zero residual solutions, the asymptotic speed of convergence of a Gauss-Newton method is that of Newton's method. This means, that if the Galerkin problem (7) has a solution, a Gauss-Newton scheme applied to (9) will find it fast and efficient (see also Section 8). Note that equations (7)-(9) represent nonlinear problems in only k variables, which may be much easier to solve than the original problem. If these smaller non-linear problems can be formulated cheaply, then the costs for an update step may be considered as being relatively small. Note also that since equations (7)-(9) are nonlinear, they may have more than one solution. This fact may be exploited to steer the computational process to a specific preferable solution of the original problem. Accelerated Inexact Newton. For the Galerkin approach, step 2e in the Inexact Newton's algorithm, Alg. 1, is replaced by four steps in which ffl the search subspace V k\Gamma1 is expanded by an approximate "Newton correction" and a suitable basis is constructed for this subspace, ffl a shadow space W k is selected on which we project the original problem, ffl the projected problem (7) is solved, ffl and the solution is updated. This is represented, by the steps 3e-3h in Alg. 3. The Minimal Residual approach and the Restricted Minimal Residual approach can be represented in a similar way. 4. Computational considerations. In this section we make some comments on implementation details that mainly focus on limiting computational work and memory space. 4.1. Restart. For small k, problems (7)-(9) are of small dimension and may often be solved at relatively low computational costs (e.g., by some variant of Newton's method). For larger k they may become a serious problem in itself. In such a situation, one may wish to restrict the subspaces V and W to subspaces of smaller dimension (see Alg. 3, step 3i). Such an approach limits the computational costs per iteration, but it may also have a negative effect on the speed of convergence. For example, the simplest choice, restricting the search subspace to a 1-dim. sub-space leads to Damped Inexact Newton methods, where, for instance, the damping parameter ff is the solution of min ff kG k (ff)k 2 , where G k ACCELERATED INEXACT NEWTON SCHEMES 7 1. choose an initial approximation u 0 . 2. 3. Repeat until u k is accurate (b) Compute the residual r (c) Compute an approximation J k for the Jacobian F 0 (u k ). (d) Solve the correction equation (approximately). Compute an (approximate) solution p k for the correction equation Expand the search space. Select a v k in the span(V that is linearly independent of V k\Gamma1 and update (f) Expand the shadow space. Select a w k that is linearly independent of W k\Gamma1 and update W (g) Solve the projected problem. Compute nontrivial solutions y of the projected system Update. Select a y k (from the set of solutions y) and update the approximation: (i) Restart. large, select an ' select ' \Theta ' 0 matrices R V and RW and compute suitable combinations of the columns of Algorithm 3: Accelerated Inexact Newton Of course, a complete restart is also feasible, say after each mth step (cf. step 3i of Alg. 3): The disadvantage of a complete restart is that we have to rebuild subspace information again. Usually it leads to a slower speed of convergence. It seems like an open door to suggest that parts of subspaces better be retained at restart, but in practical situations it is very difficult to predict what those parts should be. A meaningful choice would depend on spectral properties of the Jacobian as well as on the current approximation. When solving linear equations with GMRESR [29], good results have been reported in [30] when selecting a number of the first and the last columns (cf. step 3i of Alg. 3); e.g., 8 D. R. FOKKEMA, G. L. G. SLEIJPEN, AND H. A. VAN DER VORST In [7], a variant of GMRESR, called GCRO, is proposed, which implements another choice. For subspaces of dimension l + m, the first l columns are retained, together with a combination of the last m columns. This combination is taken such that the approximate solution, induced by a minimal residual condition, is the same for both the subspaces of dimension l +m and l + 1. To be more specific, if u where y k solves min y replaced by [V k l ; V km y km (denoting 4.2. Update. In the update step (step 3h of Alg. 3), a solution y k of the projected problem has to be selected from the set of solutions y. Selection may be necessary since many nonlinear problems have more than one solution. Sometimes, this may be the reason for poor convergence of (Inexact) Newton: the sequence of approximate solutions "wavers" between different exact solutions. For larger search subspaces, the search subspace may contain good approximations for more than one solution. This may be exploited to steer the sequence of approximate solution to the wanted solution may help to avoid wavering convergence behavior. The selection of y k should be based on additional properties of the solution u k+1 . For instance, we may look for the solution largest in norm, or as in the case of eigenvalue problems, for a solution of which one component is close to some specific value (for instance, if one is interested in eigenvalues close to, say, 0, the Ritz vector with Ritz value closest to 0 will be chosen). 4.3. The projected problem. Even though problems of small dimension can be solved with relatively low computational costs, step 3g in Alg. 3 is not necessarily inexpensive. The projected problem is embedded in the large subspace and it may require quite some computational effort to represent the problem in a small subspace (to which y belongs) of dimension ' := dim(span(V k In the case of linear equations (or linear eigenvalue problems) the computation of an ' \Theta '-matrix as W products. For this type of problems, and for many others as well, one may save on the computational costs by re-using information from previous iterations. 4.4. Expanding the search subspace. The AIN algorithm breaks down if the search subspace is not expanded. This happens when p k belongs to the span of (or, in finite precision arithmetic, if the angle between p k and this subspace is very small). Similar as for GMRES, one may then replace p k by J k v ' , where v ' is the last column vector of the matrix V k\Gamma1 . With approximate solution of the correction equation, a breakdown will also occur if the new residual r k is equal to the previous residual r k\Gamma1 . We will have such a situation if y instead of modifying the expansion process in iteration number k, one may also take measures in iteration number in order to avoid In [29] a few steps by LSQR are suggested when the linear solver is a Krylov subspace method: may already cure the stagnation. 5. How linear solvers fit in the AIN framework. In this section we will show how some well-known iterative methods for the solution of linear systems fit in the AIN framework. The methods that follow from specific choices in AIN are equivalent to well-known methods only in the sense that, at least in exact arithmetic, ACCELERATED INEXACT NEWTON SCHEMES 9 they produce the same basis vectors for the search spaces, the same approximate solutions, and the same Newton corrections (in the same sense as in which GMRES and ORTHODIR are equivalent). With the linear equation (2) is equivalent to the one in (1) and J \GammaA. In this section, M denotes a preconditioning matrix for A (i.e., for a vector v, M \Gamma1 v is easy to compute and approximates A \Gamma1 v). 5.1. GCR. With the choice, Alg. 3 (without restart) is equivalent to preconditioned GCR [31]. 5.2. FOM and GMRES. The choice algorithms that are related to FOM and GMRES [26]. With the additional choice w Alg. 3 is just FOM, while the choice gives an algorithm that is equivalent to GMRES. 5.3. GMRESR. Taking w k such that w k is perpendicular to GCR and taking p k as an approximate solution of the equation Alg. 3 is equivalent to the GMRESR algorithms [29]. One might compute p k by a few steps of GMRES, for instance. 6. AIN schemes for mildly nonlinear problems. In this section we will discuss numerical methods for the iterative solution of the generalized eigenproblem (3). We will show that they also fit in the general framework of the AIN Alg. 3 methods. As already mentioned these AIN methods consist of two parts. In one part an approximate solution of the correction equation (cf. step 3d in Alg. 3) is used to extend the search space. In the other part a solution of the projected problem (cf. step 3g in Alg. 3) is used to construct an update for the approximate solution. We will start with the derivation of a more suitable form for the (Newton) correction equation for the generalized eigenproblem. After that, we will make some comments on how to solve the projected problem. The correction equation. In order to avoid some of the complications that go with complex differentiation, we will mainly focus on the numerical computation of eigenvectors with a fixed component in some given direction (rather then on the computation of eigenvectors with a fixed norm). First, let ~ u be a fixed vector with a nontrivial component in the direction of the desired eigenvector x. We want to compute approximations u k for x with a normalized component in the ~ u-direction: (u; ~ We will select # k such that the residual r k := is orthogonal to w, where w is another fixed nontrivial vector, i.e., the approximate eigenvalue # k is given by # k := w Au k =w Bu k . Consider the map F given by F (u) := and u belongs to the hyper-plane fy 2 C 1g. The Jacobian J then given by D. R. FOKKEMA, G. L. G. SLEIJPEN, AND H. A. VAN DER VORST and the correction equation reads as u; such that equation is equivalent to ~ u 0 \Gammar k# that is, p is the solution of (10) if and only if p is the solution of (11). The projected problem. For the Generalized eigenvalue problem we are in the fortunate position that all the solutions of problems of moderate size can be computed by standard methods such as for instance the QZ [19] method. However, before we can apply these methods we have to reformulate the projected problem because of the exceptional position of u k in W F The key to this reformulation is the observation that in the methods we consider the affine subspace u k +span(V k ) is equal to V k because V k contains u k by itself. Now, as an alternative to step 3g in Alg. 3, we may also compute all the solutions y of This problem can now be solved by for instance the QZ method, and after selecting y k a new approximation u k+1 is given by u 6.1. Arnoldi's method. We consider the simplified case where I , i.e., the standard eigenproblem. If we do only one step of a Krylov subspace method (Krylov dimension 1) for the solution of the correction equation (10), then we obtain for the correction Hence, . Note that this may be a poor (very) approximation, because, in general, r k 6? ~ u. The approximate eigenvector u k belongs to the search subspace span(V expanding the search subspace by the component of p k orthogonal to span(V k\Gamma1 ) is equivalent to expanding this space with the orthogonal component of Au k , which would be the "expansion" vector in Arnoldi's method. Hence, the search subspace is precisely the Krylov subspace generated by A and u 0 . Ap- parently, Arnoldi's method is an AIN method (with a ``very inexact Newton step'') without restart. The choice W corresponds to the standard one in Arnoldi and produces #'s that are called Ritz values, while the choice W leads to Harmonic Ritz values [22]. 6.2. Davidson's method. As in Arnoldi's method, Davidson's method [6] also carries out only one step of a Krylov subspace method for the solution of the correction equation. However, in contrast to Arnoldi's method, Davidson also incorporates a preconditioner. ACCELERATED INEXACT NEWTON SCHEMES 11 He suggests to solve (10) approximately by p k with where M is the inverse of the diagonal of A \Gamma # k B. Other choices have been suggested as well (cf. e.g., [5, 21]). Because of the preconditioner, even if I , the search space is not simply the Krylov subspace generated by A and u 0 . This may lead to an advantage of Davidson's method over Arnoldi's method. For none of the choices of the preconditioner, proper care has been taken of the projections (see (10)): the preconditioner should approximate the inverse of the projected matrix (see (10)) as a map from ~ rather than of A \Gamma # k B. However, if M is the diagonal of A \Gamma #B, and we choose ~ u and w equal to the same, arbitrary standard basis vector (as Davidson does [6]) then . Note that p ? because M is diagonal and r k ? w. Therefore, for this particular choice of w (and ~ u), the diagonal M may be expected to be a good preconditioner for the correction equation (including the projections) in the cases where M is a good preconditioner for A \Gamma # k B. Observe that this argument does not hold for non-diagonal preconditioners M . 6.3. Jacobi-Davidson. Davidson methods with a non-diagonal preconditioner do not take care properly of the projections in the correction equation (10). This observation was made in [28], and a new algorithm was proposed for eigenproblems by including the projections in the Davidson scheme. In addition, these modified schemes allow for more general approximate solutions p k than . For instance, the use of ' steps of a preconditioned Krylov subspace method for the correction equation is suggested, leading to Arnoldi type of methods in which the variable polynomial pre-conditioning is determined efficiently and the projections have been included correctly. The new methods have been called Jacobi-Davidson methods (Jacobi took proper care of the projections, but did not build a search subspace as Davidson did (see [28] for details and further references)). The analysis and results in [3, 27] show that these Jacobi-Davidson methods can also be effective for solving generalized eigenproblems, even without any matrix inversion The Jacobi-Davidson methods allow for a variety of choices that may improve efficiency of the steps and speed of convergence and are good examples of AIN methods in which the projected problem (7) is used to steer the computation. For an extensive discussion, we refer to [27]. 7. AIN schemes for general nonlinear problems. In this section we summarize some iterative methods for the solution of nonlinear problems that have been proposed by different authors, and we show how these methods fit in the AIN frame-work Brown and Saad [4] describe a family of methods for solving nonlinear problems. They refer to these methods as nonlinear Krylov subspace projection methods. Their 12 D. R. FOKKEMA, G. L. G. SLEIJPEN, AND H. A. VAN DER VORST modifications to Newton's method are intended to enhance robustness and are heavily influenced by ideas presented in [9]. One of their methods is a variant of Damped Inexact Newton, in which they approximate the solution of the correction equation by a few steps of Arnoldi or GMRES and determine the damping parameter ff by a "linesearch backtracking technique". So this is just another AIN scheme, with a special 1-dimensional subspace acceleration. They also propose a model trust region approach, where they take their update to the approximation from the Krylov subspace e Vm generated by m steps of (preconditioned) Arnoldi or GMRES as Vm ~ ~ y k is the point on the dogleg curve for which k~y k k the trust region size: ~ y k is an approximation for min y Vm ~ This could be considered as a block version of the previous method. In [2] Axelsson and Chronopoulos propose two nonlinear versions of a (truncated) Generalized Conjugate Gradient type of method. Both methods fit in the AIN frame- work. The first method, NGCG, is a Minimal Residual AIN method with and V k orthonormal; in other words the correction equation is not solved. The second method, NNGCG, differs from NGCG in that p k is now computed as an approximate solution (by some method) of the correction equation (6), where the accuracy is such that non-increasing sequence [8]). So the method NNGCG is a Minimal Residual AIN method. It can be viewed as generalization of GMRESR [29]. Under certain conditions on the map F they prove global convergence. In [15], Kaporin and Axelsson propose a class of nonlinear equation solvers, GNKS, in which the ideas presented in [4] and [2] are combined. There, the direction vectors are obtained as linear combinations of the columns of e Vm and V k . To be more precise, This problem is then solved by a special Gauss-Newton iteration scheme, which avoids excessive computational work, by taking into account the acute angle between r k and J k p k , and the rate of convergence. The method generalizes GCRO [7]. 8. Numerical Experiments. In this section we test several AIN schemes and present results of numerical experiments on three different nonlinear problems. For tests and test results with methods for linear- and eigenproblems we refer to their references. The purpose of this presentation is to show that acceleration may be useful also in the nonlinear case. By useful, we mean that additional computational cost is compensated for by faster convergence. Different AIN schemes distinguish themselves by the way they (approximately) solve the correction equation and the projected problem (cf. Section 3.2 and 7). Out of the overwhelming variety of choices we have selected a few possible combinations, some of which lead to AIN schemes that are equivalent to already proposed methods and some of which lead to new methods. We compare the following (existing) Minimal Residual AIN schemes: ffl linesearch, the backtracking linesearch technique [4, pp. 458]; ffl dogleg, the model trust region approach as proposed in [4, pp. 462]; ffl nngcg, a variant of the method proposed in [2], solving (8) by the Levenberg-Marquardt algorithm [20]; ACCELERATED INEXACT NEWTON SCHEMES 13 ffl gnks, the method proposed in [15]; and the (new) Restricted Minimal Residual AIN schemes: ffl rmr a, choosing W ffl rmr b, choosing W is the component of J k p k orthogonal to W k\Gamma1 . For these last two schemes, the minimization problem solved by the Gauss-Newton variant described in [15]. The necessary subspaces for the direction p k or the projected problem were obtained by 10 steps of GMRES, or (in the third example) also by at most 50 iterations of the generalized CGS variant CGS2 [10]. In all cases the exact Jacobian was used. Furthermore, we used orthonormal matrices V k and W k , obtained from a modified Gram-Schmidt process and restricted to the last 10 columns in an attempt to save computational work. The computations were done on a Sun Sparc 20 in double precision and the iterations were stopped when method failed, either when the convergence was too slow, i.e., when or when to number of nonlinear iterations (per step) exceeded 200. Since the computational cost of the methods is approximately proportional to the costs of the number of function evaluations and matrix multiplications, the following counters are given in the tables: ffl ni, the number of nonlinear iterations; ffl fe, the number of function evaluations; ffl mv, the number of multiplications by the Jacobian; ffl pre, the number of applications of the preconditioner; ffl total, the sum of fe, mv and pre. 8.1. A 1D Burgers' Equation. As a first test problem we consider the following 1D Burgers' Equation [14] @t @x We discretized the spatial variable x with finite differences with 64 grid points and for the time derivative we used \Deltat with denotes the solution at time t n\Deltat. For this test the solution u n was computed for and as an initial guess to u n+1 we took u n . No preconditioning was used. In table Tab. 1 we show the results for problem (12) with 14 D. R. FOKKEMA, G. L. G. SLEIJPEN, AND H. A. VAN DER VORST Method ni fe mv total linesearch 594 1818 4853 6671 dogleg 644 3559 6140 9699 nngcg 229 4426 11769 16195 gnks 187 852 7067 7919 rmr a 230 926 4471 5397 Table 1: Results for Burgers' Equation A plot of the solutions is given in Fig. 1. The table shows the cumulative value of the counters for each method after completing the computation of u If we look at the number of nonlinear iterations (ni), we see that acceleration indeed reduces this number. However, in the case of gnks this does not result in less work, because the number of matrix multiplications (mv) increases too much. Here both the Galerkin approaches rmr a and rmr b are less expensive than all the other methods. rmr a being the winner. x Figure 1: Solution of Burgers' Equation 8.2. The Bratu problem. As a second test problem we consider the Bratu problem [12, 4]. We seek a solution (u; -) of the nonlinear boundary value problem: For\Omega we took the unit square and we discretized with finite differences on a 31 \Theta 31 regular grid. It is known, cf. [12], that there exist a critical value - such that for problem (13) has two solutions and for - problem (13) has no solutions. In order to locate this critical value we use the arc length continuation method as described in [12, section 2.3 and 2.4]. Problem (13) is replaced by a problem ACCELERATED INEXACT NEWTON SCHEMES 15 Method ni fe mv pre total linesearch 391 1013 3732 3421 8166 dogleg 381 2664 3010 3010 8684 nngcg 361 1297 4243 3091 8631 gnks 358 1056 6896 2780 10732 rmr a 389 539 4005 3399 7943 Table 2: Results for the Bratu Problem, solved by the arc length continuation method. Method ni fe mv pre total linesearch 29 85 336 308 729 dogleg gnks 38 119 1806 370 2295 rmr a 6 13 77 55 145 Table 3: Single solve of the Bratu Problem, u of the form where ', a scalar valued function, is chosen such that s is some arc length on the solution branch and u s is the solution of (13) for -(s). We preconditioned GMRES by ILU(0) [18] of the discretized Laplace operator \Delta. The first table Tab. 2 shows the results after a full continuation run: starting from the smallest solution (u; -) with solution branch is followed along the (discretized) arc with s we see that acceleration may be useful, in spite of the fact that there is little room for it, because on the average approximately only 4.5 Newton iterations where necessary to compute the solution per continuation step. In this example rmr b performs better than rmr a. Table Tab. 3 shows the results for the case where we solve (13) for fixed (near the critical value). In this case Galerkin acceleration is even more useful and the differences are more pronounced. The sup norm of the solution for the different values of - are plotted in Fig. 2. The two solutions at - 4 along the diagonal of the unit square are shown in Fig. 3. 8.3. The Driven Cavity Problem. In this Section we present test results for the classical driven cavity problem from incompressible fluid flow. We follow closely D. R. FOKKEMA, G. L. G. SLEIJPEN, AND H. A. VAN DER VORST Figure 2: Sup norms of the solution u along the arc. Figure 3: Solutions u at - 4 along the diagonal of the domain. the presentations in [12, 4]. In stream function-vorticity formulation the equations are @/ @n where\Omega is the unit square and the viscosity - is the reciprocal of the Reynolds number Re. In terms of / alone this can be written as subject to the same boundary conditions. This equation was discretized with finite differences on a 25 \Theta 25 grid, see Fig. 4. The grid lines are distributed as the roots of the Chebychev polynomial of degree 25. As preconditioner we used the Modified [13] decomposition of the biharmonic operator . Starting from the solution for computed several solutions, using the the arc length continuation method (cf. the previous example, and [12]) with step sizes \Deltas = 100 for 0 - Re - Tab. 4 shows the results of this test when using 10 steps of GMRES and CGS2 [10] for the correction equation. In the case of CGS2 we approximately solved the correction equation to a relative residual norm precision of 2 \Gammak , where k is the current Newton step [8], with a maximum of 50 steps. Clearly, the methods using (the basis produced steps of GMRES perform very poorly for this example. Only gnks is able to complete the full continuation run, but requires a large number of Newton steps. If we look at the results for the AIN schemes for which CGS2 is used, we see that, except for the linesearch method, these methods perform much better. The Restricted Minimal Residual methods are again the most efficient ones. ACCELERATED INEXACT NEWTON SCHEMES 17 GMRES Method ni fe mv pre total linesearch fails at Re = 400 after a total of 545. dogleg fails at Re = 100 after a total of 113. nngcg fails at Re = 2200 after a total of 19875 gnks 641 2315 30206 6210 38731 rmr a fails at Re = 2000 after a total of 13078 rmr b fails at Re = 800 after a total of 7728 Method ni fe mv pre total linesearch fails at Re = 1300 after a total of 4342. rmr a 137 297 6266 5969 12532 Table 4: Results for the Driven Cavity problem, solved by the arc length continuation method for Re Figure 4: Grid for the Driven Cavity problem, (25 \Theta 25). Figure 5: Stream lines of the Driven Cavity problem, Re = 100. This test also reveals a possible practical drawback of methods like dogleg and gnks. These methods exploit an affine subspace to find a suitable update for the approximation. This may fail, when the problem is hard or when the preconditioner is not good enough. In that case the dimension of the affine subspace must be large, which may be, because of storage requirements and computational overhead, not fea- sible. For the schemes that use approximate solutions of the correction equation, delivered by some arbitrary, iterative method, e.g., CGS2, one can easily adapt the precision, which leaves more freedom. D. R. FOKKEMA, G. L. G. SLEIJPEN, AND H. A. VAN DER VORST Figure lines of the Driven Cavity problem, Re = 400. Figure 7: Stream lines of the Driven Cavity problem, Re = 1600. Figure 8: Stream lines of the Driven Cavity problem, Re = 2000. Figure 9: Stream lines of the Driven Cavity problem, Re = 3000. Plots of the stream lines for the values 0:0; 0:0025; 0:001; 0:0005; 0:0001; 0:00005 (cf. [12]) are given in Fig. 5-9. The plots show virtually the same solutions as in [12]. 9. Conclusions. We have shown how the classical Newton iteration scheme for nonlinear problems can be accelerated in a similar way as standard Richardson-type iteration schemes for linear equations. This leads to the AIN framework in which ACCELERATED INEXACT NEWTON SCHEMES 19 many well-known iterative methods for linear-, eigen-, and general nonlinear problems fit. From this framework an overwhelming number of possible iterations schemes can be formulated. We have selected a few and shown by numerical experiments that especially the Restricted Minimal Residual methods can be very useful for further reducing computational costs. --R The principle of minimized iterations in the solution of the matrix eigenvalue problem On nonlinear generalized conjugate gradient methods te Riele Hybrid Krylov methods for nonlinear systems of equations The Davidson method The iterative calculation of a few of the lowest eigenvalues and corresponding eigenvectors of large real symmetric matrices Nested Krylov methods and preserving the orthogonality Generalized conjugate gradient squared QMR: A quasi minimal residual method for non-Hermitian linear systems A class of first order factorizations methods The partial differential equation u t On a class of nonlinear equation solvers based on the residual norm reduction over a sequence of affine subspaces Acceleration techniques for decoupling algorithms in semiconductor simulation Solution of systems of linear equations by minimized iteration An iterative solution method for linear systems of which the coefficient matrix is a symmetric M-matrix An algorithm for generalized matrix eigenvalue problems Generalizations of Davidson's method for computing eigenvalues of large non-symmetric matrices Approximate solutions and eigenvalue bounds from Krylov subspaces Krylov subspace method for solving large unsymmetric linear systems GMRES: A generalized minimum residual algorithm for solving nonsymmetric linear systems A Jacobi-Davidson iteration method for linear eigenvalue problems GMRESR: A family of nested GMRES methods Further experiences with GMRESR Generalized conjugate-gradient acceleration of nonsymmetrizable iterative methods --TR --CTR Keith Miller, Nonlinear Krylov and moving nodes in the method of lines, Journal of Computational and Applied Mathematics, v.183 n.2, p.275-287, 15 November 2005 P. R. Graves-Morris, BiCGStab, VPAStab and an adaptation to mildly nonlinear systems, Journal of Computational and Applied Mathematics, v.201 n.1, p.284-299, April, 2007 Heng-Bin An , Ze-Yao Mo , Xing-Ping Liu, A choice of forcing terms in inexact Newton method, Journal of Computational and Applied Mathematics, v.200 n.1, p.47-60, March, 2007 Heng-Bin An , Zhong-Zhi Bai, A globally convergent Newton-GMRES method for large sparse systems of nonlinear equations, Applied Numerical Mathematics, v.57 n.3, p.235-252, March, 2007 Monica Bianchini , Stefano Fanelli , Marco Gori, Optimal Algorithms for Well-Conditioned Nonlinear Systems of Equations, IEEE Transactions on Computers, v.50 n.7, p.689-698, July 2001

nonlinear problems;iterative methods;newton's method;inexact Newton

289882

Quantitative Evaluation of Register Pressure on Software Pipelined Loops.

Software Pipelining is a loop scheduling technique that extracts loop parallelism by overlapping the execution of several consecutive iterations. One of the drawbacks of software pipelining is its high register requirements, which increase with the number of functional units and their degree of pipelining. This paper analyzes the register requirements of software pipelined loops. It also evaluates the effects on performance of the addition of spill code. Spill code is needed when the number of registers required by the software pipelined loop is larger than the number of registers of the target machine. This spill code increases memory traffic and can reduce performance. Finally, compilers can apply transformations in order to reduce the number of memory accesses and increase functional unit utilization. The paper also evaluates the negative effect on register requirements that some of these transformations might produce on loops.

Introduction Current high-performance floating-point microprocessors try to maximize the exploitable parallelism by: heavily pipelining functional units (1;2) , making aggressive use of parallelism (3;4) , or a combination of both (5) which is the trend in current and future high-performance microprocessors. To exploit effectively this amount of available parallelism, aggressive scheduling techniques such as software pipelining are required. Software pipelining (6) is an instruction scheduling technique that exploits the Instruction Level Parallelism (ILP) of loops by overlapping the execution of successive iterations. There are different approaches to generate a software pipelined schedule for a loop (7) . Modulo scheduling is a class of software pipelining algorithms that relies on generating a schedule for an iteration of the loop such that when this same schedule is repeated at regular intervals, no dependence is violated and no resource usage conflict arises. Modulo scheduling was proposed at the beginning of the 80s (8) . Since then, many research papers have appeared on the topic (9;10;11;12;13;14;15;16) , and it has been incorporated into some production compilers (17;18) . The drawback of aggressive scheduling techniques such as software pipelining is that they increase register requirements. In addition, increasing either the stages of functional units or the number of functional units, which are the current trends in microprocessor design, tends to increase the number of registers required by software pipelined loops (19) . The register requirements of a schedule are of extreme importance for compilers since any valid schedule must fit in the available number of registers of the target ma- chine. In this way there has been a research effort to produce optimal/near-optimal modulo schedules with minimum/reduced register requirements. The optimal methods are mainly based on linear programming approaches (20;21) . Unfortunately, optimal techniques have a prohibitive computational cost which make them impractical for product compilers. Some practical modulo scheduling approaches use heuristics to produce near-optimal schedules with reduced register requirements (11;16;22) . Other approaches try to reduce the register requirements of the schedules by applying a post-pass process (23;24) . There have been proposals to perform register allocation for software pipelined loops (25;26;27) . If the number of registers required is larger than the available number of registers, spill code has to be introduced to reduce register usage. Different alternatives to generate spill code for software pipelined loops have been proposed and evaluated in (28) . This spill code can also reduce performance. In this paper we show that the performance and memory traffic of aggressive (in terms of ILP) machines are heavily degraded due to a lack of registers. To avoid this performance degradation big register files are required. In addition, the number of read and write ports increases with the number of functional units. The number of registers and the number of ports have a negative effect on the area required by the register file (29) and on the access time to the register file (30) . Some new register file organizations have been proposed to have big register files (in terms of registers and access ports) without degrading either area or access time (31;32;33;34) . Some of these organizations have been used, combined with register-sensitive software pipelinig techniques (35) , resulting in better performance. This shows that most of the techniques for reducing the effects of register pressure are complementary. In this paper, register requirements of pipelined floating-point intensive loops are evaluated. Several studies are performed. First, the register requirements of loop invariants, which are a machine-independent characteristic of the loops, are studied in Section 3. Section 4 carries out a study of the register requirements of loop variants as a function of the latency and the number of functional units. Section 5 studies the cumulative register requirements of both (loop variants and invariants) and show that loops with high register requirements represent a high percentage of the execution time of the Perfect Club. Section 6 considers the effects of a limited size register file on performance and memory traffic; for this purpose spill code has been added to those loops that require more registers than are available. Finally, Section 7 analyzes the effects on the register requirements of some optimizations that try to improve performance by reusing data and increasing functional unit usage. 2. Related Concepts and Experimental Framework 2.1. Overview of Modulo Scheduling In a software pipelined loop, the schedule for an iteration is divided into stages so that the execution of consecutive iterations that are in distinct stages is overlapped. The number of stages in one iteration is termed stage count (SC). The number of cycles per stage is the initiation interval (II). The execution of a loop can be divided into three phases: a ramp up phase that fills the software pipeline, a steady state phase where the software pipeline achieves maximum overlap of iterations, and a ramp down phase that drains the software pipeline. During the steady state phase of the execution, the same pattern of operations is executed in each stage. This is achieved by iterating on a piece of code, termed the kernel, that corresponds to one stage of the steady state phase. The initiation interval II between two successive iterations is bounded either by loop-carried dependences in the graph (RecMII) or by resource constraints of the architecture (ResMII). This lower bound on the II is termed the Minimum Initiation Interval (MII= max (RecMII, ResMII)). The reader is referred to (18;13) for an extensive dissertation on how to calculate ResMII and RecMII. 2.2. Register Requirements Values used in a loop correspond either to loop-invariant variables or to loop- variant variables. Loop invariants are repeatedly used but never defined during loop execution. Loop invariants have a single value for all the iterations of the loop; each invariant requires one register regardless of the scheduling and the machine configuration. For loop variants, a value is generated in each iteration of the loop and, therefore, there is a different value corresponding to each iteration. Because of the nature of software pipelining, lifetimes of values defined in an iteration can overlap with lifetimes of values defined in subsequent iterations. Lifetimes of loop variants can be measured in different ways depending on the execution model of the machine. We assume that a variable is alive from the beginning of the producer operation, until the start of the last consumer operation. By overlapping the lifetimes of the different iterations, a pattern of length II cycles that is indefinitely repeated is obtained. This pattern indicates the number of values that are live at any given cycle. As it is shown in (26) , the maximum number of simultaneously live values (MaxLive) is an accurate approximation of the number of registers required by the schedule. Values with a lifetime greater than II pose an additional difficulty since new values are generated before previous ones are used. One approach to fix this problem is to provide some form of register renaming so that successive definitions of a value use distinct registers. Renaming can be performed at compile time by using modulo variable expansion (MVE) (36) , i.e., unrolling the kernel and renaming at compile time the multiple definitions of each variable that exist in the unrolled kernel. A rotating register file can be used to solve this problem without replicating code by renaming different instantiations of a loop variant at execution time (37) . In this paper we assume the presence of rotating register files. 2.3. Experimental Framework The experimental evaluation has been done using all the innermost loops of the Perfect Club Benchmark Suite (38) that are suitable for software pipelining. These loops have been obtained with the ICTINEO compiler (39) . ICTINEO is a source-to- source restructurer developed on top of Polaris (40) with an internal representation that combines both high and low-level information. It includes some basic transformations that allow us to obtain optimized data dependence graphs. A total of 1258 loops suitable for software pipelining have been used. This set includes all the innermost loops that do not have subroutine calls or conditional exits. Loops with conditional structures in their bodies have been IF-converted (41) , with the result that the loop now looks like a single basic block. The loops represent 78% of the total execution time of the Perfect Club measured on a HP-PA 7100. In addition, to show the effects of aggressive optimizations in individual loops, we have also used some Livermore Loops (42) . The loops have been scheduled for a wide range of VLIW-like target configura- tions, with different number of functional units and latencies. Table 1 shows the different configurations used along the paper. All functional units are fully pipelined, except the divider, which is not pipelined at all. We labeled the different configurations by PxLy where x is the number of functional units of each kind and y is the latency (number of stages) of the mostly used functional units, that is, the adder and the multiplier. We considered a constant latency for loads, stores, divisions and square roots independently of the configuration. Different metrics are used along the paper to evaluate performance. On the one side, the register requirements are evaluated computing the maximum number of simultaneously live values MaxLive. It has been shown (26) that some allocation strategies almost always achieve the MaxLive lower bound. In particular, the wands- only strategy using end-fit with adjacency ordering almost never requires more than MaxLive registers. Second, execution times are approximated as the II obtained after modulo scheduling times the trip count of the loop. Finally, memory traffic is approximated evaluating the number of memory accesses that are needed to execute a loop, considering both the memory accesses defined in the graph and the spill code introduced due to having a finite register file. 3. Register Requirements of Loop Invariants Loop invariants are values that are repeatedly used by a loop at each iteration, but never written by it. That is, they are defined before entering the loop, are used by it, and are not redefined at least until the loop has finished. Loop-invariant variables can be either stored in registers or in memory but, since memory bandwidth is without doubt the most performance limiting factor in current processors, even the most simple optimizing compilers try to hold them in registers during the execution of the loop. For this purpose loop-invariant variables are loaded from memory to a register before entering the loop that uses them. Even assuming no memory bandwidth restrictions loading these variables to registers before entering the loop saves instruction bandwidth in load/store architectures, i.e. almost all current processors. Another source of loop invariants are the invariant computations (i.e. computations that produce the same result during all the iterations of the loops). These computations can be extracted out of the loop -where they are performed at every iteration- and computed only once before entering the loop. Also because of memory bandwidth and instruction bandwidth, the partial result of these computations is held in registers. This optimization is termed loop-invariant removal and is one of the optimizations performed by the ICTINEO compiler. In fact, if we consider load and store operations as computations, they also can be extracted during the loop-invariant removal optimization. Figure 1a shows an example. Figure 1b shows a low-level representation of the loops before removing loop-invariant computations. Figure 1c shows the same loops after removing loop-invariant computations. Notice that q and are loop-invariant variables for both loops, so they can be loaded to registers before entering them. C(j) is also a loop-invariant variable for the innermost loop, so it can be extracted from it, but not from the outermost loop. Because v1 and v2 are loop invariants (which are assumed to be allocated in registers), the computation v3 is a loop-invariant computation, and therefore can be extracted from the innermost loop. The extraction of loop invariants has lead to a smaller innermost loop -the one that is executed most of the time- with less operations. In the original one there were: 1 addition, 2 multiplications and 5 memory accesses, while in the optimized one there are only 1 addition, 1 multiplication and 2 memory accesses. The ICTINEO compiler performs -among other optimizations- an aggressive extraction of loop-invariant computations. We have used the optimized dependence graph to evaluate the register requirements of loops due to loop invariants. Figure 2 shows the cumulative distribution of the requirements for loop invariants for all 1258 loops. In general loops have very few loop invariants. For instance 25% of the loops have no loop invariants and 95% of the loops have 8 or less invariants. Nevertheless a few loops have a high number of loop invariants. For instance 9 loops have more than invariants and 1 of them requires 68 loop invariants. 4. Register Requirements of Loop Variants Unlike loop invariants, the number of registers required by loop variants is a schedule dependent characteristic. So, the register requirements depend on the scheduling technique, and the target machine configuration for which the scheduling is performed, as well as the topology of the loop. In this section we are interested in the effects of the machine configuration on the register requirements. The main characteristics that influence the final schedule, and therefore the register requirements, are the number of stages of functional units and the number of functional units. That is, the degree of pipelining and the degree of parallelism. For this purpose we have generated the software pipelined schedule for all the Perfect Club loops for the target configurations shown in Table 1. Figure 3 shows the cumulative distribution of registers for each configuration. Notice that when the product of the latency of functional units and the number of functional units goes up, the number of registers needed by the loops also increases. Note that the number of functional units has a slightly bigger effect on the number of registers required than the degree of pipelining has. This is mainly due to the fact that the latency of some functional units is kept constant for all the configurations. There is a small number of loops (6%) that do not require any register for loop variants. The loop bodies of these loops have no invariants because they only have store operations. These loops are typically used to initialize data structures. There is another region (6% to 30%) where the loops have few register requirements and where the register requirements of the loops are not influenced by the number of stages. This is due to the fact that these loops have no arithmetic opera- tions, that is, their bodies only have load and store operations. These loops have few register requirements to hold the values between loads and stores. They are basically used to copy data structures. 5. Combined Register Requirements Although some special architectures, such as the Cydra-5 (43) , have separated register files for loop-invariant variables (global register file) and loop-variant variables (rotating register file), all the current superscalar microprocessors, for which software pipelining can also be applied, have a single register file to store both sets of variables. Therefore it is of great interest to know the combined register pressure of loop-variant variables plus loop-invariant variables. Figure 4 shows the cumulative distribution of the register requirements of the loops of the Perfect Club when software pipelined for the configurations of Table 1. It is interesting to notice that 96% of the loops can be scheduled with registers and without adding spill code for the less aggressive configuration (P1L2). On the contrary only 85% of the loops can be scheduled with registers for the more aggressive configuration (P2L6). If 64 registers were available we would be able to schedule 99.5% and 96% of the loops respectively. Also, approximately 60% of the loops (it varies depending on the configuration) require a small number of registers or less). From these figures, one can conclude that register requirements are not extremely high, and that 64 registers might seem enough for the configurations used. Even though, we have observed that small loops represent a small percentage of the execution time and that most of the time is spent in big loops which, in general, have higher register requirements. Figure 5 shows the dynamic register requirements, where each loop has been weighted by its execution time. The graph of Figure 5 is similar to the one of Figure 4 but instead of representing the percentage of loops that require a certain amount of registers, it represents the percentage of time spent on loops that require a certain amount of registers. Data gathered from this figure shows that small loops requiring less that 8 registers portion of the execution time (about 15%). If only 32 registers are available, the loops that can be scheduled without adding spill code represent only between 67% and 52% of the execution time. And even for a machine with 64 registers, the loops that can be scheduled without adding spill code represent only between 78% and 69% of the execution time. Even more, a big percentage of the execution time of the Perfect Club is spent on a few loops that require more than 100 registers. 6. Effects of a Limited Register File In the previous sections an infinite number of registers has been assumed. In this section we study the effects of having a limited amount of registers on performance. When there is a limited number of registers and the register allocator fails to find a solution with the number of registers available, some additional action must be taken. In (28) we have proposed several alternatives to schedule software pipelined loops with register constraints. For the purposes of our experiments, we use the best option, in terms of performance for the loops, to add spill code. Current microprocessors have only floating-point registers and 32 integer reg- isters. We think that future generations of microprocessors will enlarge the register file to 64 registers (at least for the floating point-register file). Since this study is targeted to floating point intensive applications, we study the effects of having register files with registers and with 64 registers compared to the ideal case of having an infinite number of registers. Adding spill code to generate a valid schedule with a given number of registers produces two negative effects. One of them can hurt performance indirectly, since spill code adds new load and store operations, which might interfere in the memory subsystem or cause additional cache misses. The other effect, that affects performance directly, is that if new operations are added it might be necessary to increase the II of the loop, reducing the throughput even in the hypothetical case of having a perfect memory system. Figure 6 shows the number of memory accesses required to execute all the loops for the six configurations we use. Notice that the number of memory accesses is the same for all the configurations if no spill code is added. Also predictable is the fact that the number of memory accesses increases as the number of registers is reduced. It can also be observed that the growth of memory accesses is more dramatic for aggressive configurations. For instance the configuration P2L6 with 64 registers requires 52% more accesses than an ideal machine with an infinite register file and 176% more memory accesses if it only has registers. Anyway it is difficult to predict the performance degradation that these additional accesses can have without simulating the memory subsystem, which is out of the scope of this paper. In any case, we can easily predict the direct effect on performance that the spill code has on the execution time of the loops because of larger IIs. Figure 7 shows the number of cycles required to execute all the loops for the six configurations, with an infinite number of registers, with 64 registers, and with registers. Notice that, as expected, fewer registers mean lower performance (more cycles are required to execute all the loops). In general, the more aggressive the machine configuration is, the bigger the performance degradation. For instance, a P1L6 machine can have a speed-up of 1.19 if the register file is doubled from 32 registers to 64 registers. Instead, a P2L6 machine will have a speed-up of 1.29 by doubling the register file. Through the data gathered from this experiment it can be concluded that simply adding functional units without caring about the number of registers results in performance figures lower than expected (due to the negative effects of the additonal spill code). For instance, doubling the number of functional units (i.e. going from P1L6 to P2L6) produces a speed-up of 1.59 if the number of registers is not limited. For machines with 64 and 32 registers, the speed-up is 1.56 and 1.43, respectively. However, if the number of register is doubled together with the number of functional units (i.e. from P1L6 with 32 registers to P2L6 with 64 registers), the speed-up is 1.85. 7. Optimizations and Register Requirements In addition to the effect of latency and the number of functional units, register requirements of loops can also increase when certain optimizations are applied to the loops. In order to see how some 'advanced' optimizations affect the register requirements of loops we have hand-optimized some of the Livermore Loops and a few loops from the Perfect Club, where the optimizations are applicable. We have not considered common optimizations such as loop-invariant removal, common subexpression elimination, redundant load and store removal, or dead code removal which we assume are applied by any compiler. In the following subsections, we present a brief description of the optimizations studied and the effect that applying the optimizations has on both performance and register requirements. 7.1. Loop Unrolling Loop unrolling (44) is an optimization used by the compilers of current micro- processors. It allows a better usage of resources because several iterations can be scheduled together, increasing the number of instructions available to the scheduler. It also reduces the loop overhead caused by branching and index update. In our case we achieve efficient iteration overlapping through Software Pipelining. Unrolling is required to match the number of resources required by the loop with the resources of the processor and also to schedule loops with fractional MII (45) . As an example, assume a loop with 5 additions and a processor with 2 adders. If there are no recurrences and additional resources do not limit the scheduling, this loop can be scheduled with cycles. The loop will be executed at 83.3% of the peak machine performance. If the loop is unrolled once before scheduling, we obtain a new loop with 10 additions. This new loop can be scheduled with an but each iteration of the new loop corresponds to two iterations of the original loop, so on average an iteration is completed in 2.5 cycles. The unrolled loop can be executed at 100% of the peak machine performance. Unroll increases the register requirements because bigger loops have more temporal variables to store, which usually requires more registers. Table 2 shows the effects of unrolling some loops on their register requirements. For the loops used, unrolling them once results in a better usage of the available resources. Unfortunately it also produces an increase in the register requirements which, as we have seen in Section 6, can degrade performance if the number of available registers is less than the registers required. 7.2. Common Subexpression Elimination Across Iterations Common subexpression elimination across iterations (CSEAI) is an extension for inner loops of the common subexpression elimination optimization applied by almost all optimizing compilers. It consists of the reuse of values generated in previous iterations. The reused data has to be stored in a register and, since the value is used across iterations, more than one physical register is required in order not to overwrite a live value with a new instantiation of the same variable. This is an extension to loops of the 'common subexpression elimination' optimization. It is more sophisticated in the sense that it can only be applied to loops and that it requires a dependence analysis to know which values of the current iteration are used in further iterations. The objective of this optimization will be to reduce the number of operations of the loop body which can lead to a lower II. Even if after applying the optimization the final II is not lowered, it is worthwhile applying it if the number of memory accesses is reduced. As an example consider the loop of Figure 8a whose unoptimized body is shown in Figure 8b. We have added to each of the loop variants generated by the operations a subindex associated to the iteration where it is generated. For instance represents the value generated by the first operation at iteration i. A conventional common subexpression elimination step will recognize that both load the same value Z(i+1). Therefore, operation V 4 i be removed and all subsequent uses of V 4 i be substituted by uses of V 2 i . This optimization produces as output the loop shown in Figure 8c. Notice that because of this optimization the loop requires only 7 operations instead of 8. The common subexpression elimination optimization can detect that the value of V 3 at the previous iteration (V 3 is equivalent to the value of . Therefore the operations required to calculate V 6 i can be removed, and all uses of substituted by V 3 . The same can be done with operations producing the loop body of Figure 8d. Notice that the resulting loop body has only 4 operations. Table 3 shows the II and the memory traffic of some loops with and without applying CSEAI. Notice that in all cases there are improvements in memory traffic, in the II or in both. It is also interesting to notice that, even though the optimization reduces the number of operations of the loops and therefore the number of lifetimes per iteration, the register requirements increase in most cases. This is mainly due to the fact that some loop variants must be preserved across several iterations. 7.3. Back-substitution Back-substitution (37) is a technique that increases the parallelism of loops that have single recurrences and therefore limited parallelism, but it increases the number of operations executed per iteration. For instance, consider the Livermore Loop 11 shown in Figure 9a. There is a recurrence of weight 1 that limits the MII . If we have an adder with 6 cycles of latency, this loop can be scheduled with an II of 6 cycles. Using back-substitution once, the loop becomes the one shown in Figure 9b. Now the recurrence that limits the schedule of the loop has a weight of 2. Obviously in this case we have to do two additions instead of one, but this transformed loop can be executed with an II of 3 cycles doubling the performance. To avoid part of this increase in operations, other optimizations can be applied such as common subexpression elimination across iterations. In this example it reduces the number of loads, but not the number of additions. A way of further reducing the increment of operations is to perform unroll as well as back-substitution. If we apply unroll to Livermore Loop 11 together with back-substitution we can obtain the loop of Figure 9c. Notice that in this loop we require 3 additions to compute two iterations of the original loop. The loop of Figure 9c can be executed with an II of 6 cycles but it performs two iterations, so the performance is the same as in the loop of Figure 9b and requires fewer operations. When back-substitution is applied the number of operations per iteration is in- creased, and the II is reduced. In general, bigger loops require more registers because they have more temporal values to store; in addition a reduced II requires more registers for the same temporal values, because new values are created in each iteration. Table 4 shows the effect on both performance and register requirements when applying back-substitution to Livermore Loops 11 and 5. Notice that for loop 5 in some cases the II is 13. This is because, in those cases, the scheduler used fails to find a schedule in 12 cycles, even though it exists. 7.4. Blocking Blocking at the register level is a well-known optimization in the context of dense matrix linear algebra (46) . Blocking is a transformation applied to multiple nested loops that finds opportunities for reuse of subscripted variables and replaces the memory references involved by references to temporary scalar variables allocated in registers. Blocking basically consists of unrolling the outer loops and restructuring the body of the innermost loop so that the memory references that are reused in the same iteration are held in registers for further uses. Blocking reduces the number of memory accesses, so if the loop is memory-bound, a reduction of the number of memory references improves the maximum performance. Also, a reduction of the memory references can improve performance due to less interference in the memory subsystem. Finally, if the innermost loop is bound by recurrences, doing several iterations of the outermost loops, in one iteration of the innermost loop, can improve the resource usage. Unfortunately, blocking increases the size of the loop, which in general increases the register requirements. It also enlarges the lifetimes of some loop variants, which must be stored in registers for a longer time because they are reused later. This enlargement of lifetimes also contributes to an increase of the register requirements of the transformed loop versus the original one. Table 5 shows the II , the effective II per iteration, the memory traffic and the registers required when blocking is applied to a basic matrix by matrix kernel. 8. Conclusions In this paper we have evaluated the register requirements of software pipelined loops. We have evaluated the register requirements of loop invariants and loop variants of the loops of the Perfect Club. We have shown empirically that the register requirements of loop variants increases with the latency and the number of functional units. These results corroborate the theoretical study done in (19) . We have also shown that loops with high register requirements take up an important proportion of the execution time of some representative numerical applications. We have also evaluated the effect of register file size on the number of memory accesses and on the performance. We have shown that the memory traffic can have a high growth if the register file is small and the machine configuration is very ag- gressive. Also performance (even under the hypothesis of a perfect memory system) is degraded with small register files. Finally, we have done a limited evaluation of the effects of some advanced opti- mizations. We have shown that these advanced optimizations increase performance and reduce memory traffic at the expense of an increase in the register requirements. This suggests that, in practice, the optimizations must be performed carefully. If a loop is excessively optimized, the high register requirements can offset the benefits of the optimization applied, and even produce worse results. As future work, we will integrate those advanced optimizations in the ICTINEO compiler in order to perform an extensive evaluation of the effect of those optimizations on performance and the tradeoffs with the performance degradation due to the higher register requirements. Acknowledgements This work has been supported by the Ministry of Education of Spain under contract TIC 429/95, and by CEPBA (European Center for Parallelism of Barcelona). --R The Mips R4000 processor. Next generation of the RISC System/6000 family. Superscalar instruction execution in the 21164 Alpha microprocessor. An approach to scientific array processing: The architectural design of the AP120B/FPS-164 family Software pipelining. Some scheduling techniques and an easily schedulable horizontal architecture for high performance scientific computing. Software pipelining: An effective scheduling technique for VLIW ma- chines Circular scheduling: A new technique to perform software pipelining. Parallelisation of loops with exits on pipelined architectures. Iterative modulo scheduling: An algorithm for software pipelining loops. A realistic resource-constrained software pipelining algo- rithm Modulo scheduling with multiple initiation inter- vals Hypernode reduction modulo scheduling. Software pipelining in PA-RISC compilers Compiling for the Cydra 5. Register requirements of pipelined processors. Minimal register requirements under resource-constrained software pipelining Optimum modulo schedules for minimum register requirements. Swing modulo scheduling: A lifetime Stage scheduling: A technique to reduce the register requirements of a modulo schedule. RESIS: A new methodology for register optimization in software pipelining. Register allocation using cyclic interval graphs: A new approach to an old problem. Register allocation for software pipelined loops. The meeting a new model for loop cyclic register allocation. Heuristics for register-constrained software pipelining Principles of CMOS VLSI Design: A systems Per- spective Partitioned register files for VLIWs: A preliminary analysis of tradeoffs. Using Sacks to organize register files in VLIW machines. Digital 21264 sets new standard. Reducing the Impact of Register Pressure on Software Pipelining. A Systolic Array Optimizing Compiler. Overlapped loop support in the Cydra The Perfect Club benchmarks: Effective performance evaluation of supercomputers. A uniform representation for high-level and instruction-level transformations POLARIS: The next generation in parallelizing compilers. Conversion of control dependence to data dependence. The Livermore FORTRAN kernels: A computer test of the numerical performance range. The Cydra 5 departmental super- computer: design philosophies Unrolling loops in FORTRAN. Software pipelining: A comparison and improvement. Improving register allocation for subscripted variables. --TR Principles of CMOS VLSI design: a systems perspective Software pipelining: an effective scheduling technique for VLIW machines The Cydra 5 Departmental Supercomputer Overlapped loop support in the Cydra 5 Improving register allocation for subscripted variables Parallelization of loops with exits on pipelined architectures Circular scheduling Register allocation for software pipelined loops Register requirements of pipelined processors Partitioned register files for VLIWs Lifetime-sensitive modulo scheduling Compiling for the Cydra 5 Designing the TFP Microprocessor Iterative modulo scheduling Minimizing register requirements under resource-constrained rate-optimal software pipelining Software pipelining Optimum modulo schedules for minimum register requirements Modulo scheduling with multiple initiation intervals Stage scheduling Hypernode reduction modulo scheduling Heuristics for register-constrained software pipelining Software pipelining A Systolic Array Optimizing Compiler Conversion of control dependence to data dependence The Mips R4000 Processor Superscalar Instruction Execution in the 21164 Alpha Microprocessor RESIS Using Sacks to Organize Registers in VLIW Machines Some scheduling techniques and an easily schedulable horizontal architecture for high performance scientific computing Non-Consistent Dual Register Files to Reduce Register Pressure Swing Modulo Scheduling --CTR Javier Zalamea , Josep Llosa , Eduard Ayguad , Mateo Valero, Software and hardware techniques to optimize register file utilization in VLIW architectures, International Journal of Parallel Programming, v.32 n.6, p.447-474, December 2004 David Lpez , Josep Llosa , Mateo Valero , Eduard Ayguad, Widening resources: a cost-effective technique for aggressive ILP architectures, Proceedings of the 31st annual ACM/IEEE international symposium on Microarchitecture, p.237-246, November 1998, Dallas, Texas, United States Javier Zalamea , Josep Llosa , Eduard Ayguad , Mateo Valero, Two-level hierarchical register file organization for VLIW processors, Proceedings of the 33rd annual ACM/IEEE international symposium on Microarchitecture, p.137-146, December 2000, Monterey, California, United States David Lpez , Josep Llosa , Mateo Valero , Eduard Ayguad, Cost-Conscious Strategies to Increase Performance of Numerical Programs on Aggressive VLIW Architectures, IEEE Transactions on Computers, v.50 n.10, p.1033-1051, October 2001 Javier Zalamea , Josep Llosa , Eduard Ayguad , Mateo Valero, Improved spill code generation for software pipelined loops, ACM SIGPLAN Notices, v.35 n.5, p.134-144, May 2000

SOFTWARE PIPELINING;REGISTER REQUIREMENTS;SPILL CODE;LOOP TRANSFORMATIONS;PERFORMANCE EVALUATION

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A Feasibility Decision Algorithm for Rate Monotonic and Deadline Monotonic Scheduling.

Rate monotonic and deadline monotonic scheduling are commonly used for periodic real-time task systems. This paper discusses a feasibility decision for a given real-time task system when the system is scheduled by rate monotonic and deadline monotonic scheduling. The time complexity of existing feasibility decision algorithms depends on both the number of tasks and maximum periods or deadlines when the periods and deadlines are integers. This paper presents a new necessary and sufficient condition for a given task system to be feasible and proposes a new feasibility decision algorithm based on that condition. The time complexity of this algorithm depends solely on the number of tasks. This condition can also be applied as a sufficient condition for a task system using priority inheritance protocols to be feasible with rate monotonic and deadline monotonic scheduling.

INTRODUCTION Hard real-time systems have been defined as those containing processes that have deadlines that cannot be missed [Bur89a]. Such deadlines have been termed hard: they must be met under all circumstances otherwise catastrophic system failure may result To meet hard deadlines implies constraints upon the way in which system resources are allocated at runtime. This includes both physical and logical resources. Conventionally, resource allocation is performed by scheduling algorithms whose purpose is to interleave the executions of processes in the system to achieve a pre-determined goal. For hard real-time systems the obvious goal is that no deadline is missed. One scheduling method that has been proposed for hard real-time systems is the rate monotonic algorithm [Liu73a]. This is a static priority based algorithm for periodic processes in which the priority of a process is related to its period. Whilst this algorithm has several useful properties, including a schedulability test that is sufficient and necessary [Leh89a], the constraints that it imposes on the process system are severe: processes must be periodic, independent and have deadline equal to period. Many papers have successively weakened the constraints imposed by the rate-monotonic algorithm and have provided associated schedulability tests. Reported work includes a test to allow aperiodic processes to be scheduled [Sha89a], and a test to schedule processes that synchronise using semaphores [Sha88a]. One constraint that has remained is that the deadline and period of a process must be equal. The weakening of this latter constraint would benefit the application designer by providing a more flexible process model for implementing the system. The increased flexibility is seen by observation: processes with deadline = period are expressible within a process model permitting deadline - period . For example, process systems whose timing characteristics are suitable for rate-monotonic scheduling would also be accepted by a scheduling scheme permitting deadlines and periods of a process to differ. This paper relaxes this constraint and so transforms the rate-monotonic algorithm into the deadline-monotonic algorithm. Schedulability tests are developed which guarantee the deadlines of periodic processes. This approach is then shown to be applicable for guaranteeing the deadlines for arbitrary mixtures of periodic and sporadic processes. The following sub-section gives a brief description of the symbols and terminology used in the remainder of the paper. Section 2 gives an overview of the rate-monotonic scheduling algorithm and associated schedulability tests. Section 3 introduces the deadline-monotonic scheduling algorithm. New schedulability constraints for the algorithm are developed. Section 4 outlines some previously proposed methods of guaranteeing sporadic process deadlines within the context of the rate-monotonic algorithm. The section then proposes a simpler method guaranteeing the deadlines of arbitrary mixtures of sporadic and periodic processes using the deadline-monotonic scheduling algorithm. 1.1. Notation A process is periodic if it is released for execution in a periodic manner. When this is not the case, and a maximum release frequency can be defined, the process is termed sporadic. If no such maximum can be defined the process is termed aperiodic [Aud90a]. A process is given by t i , where i identifies the process. The subscript i is defined to be the priority of that process, where priorities are unique. Priorities are assigned numerically, taken from the interval # # where 1 is the highest priority and n (the number of processes in the system) the lowest. The process t i has timing characteristics T i , C i and D i . These refer to the value of the period, computation time and deadline of t i . 2. THE RATE-MONOTONIC SCHEDULING ALGORITHM Rate-monotonic scheduling is a static priority based mechanism [Liu73a]. Priorities assigned to processes are inversely proportional to the length of the period. That is, the process with the shortest period is assigned the highest priority. Processes are executed in preemptive manner: at any time, the highest priority process with outstanding computation requirement is executed. Amongst the class of static priority scheduling schemes, it has been shown that rate-monotonic priority assignment is optimal [Liu73a]. This implies that if a given static priority scheduling algorithm can schedule a process system, the rate-monotonic algorithm is also able to schedule that process system. In the case of the rate-monotonic scheduling algorithm, optimality implies the imposition of constraints upon the process system. These include: # fixed set of processes; all processes are periodic; # all processes have deadline equal to period; # one instance of a process must be complete before subsequent instances are run; # all processes have known worst-case execution times; no synchronisation is permitted between processes; # all processes have an initial release at time 0. The last of these constraints is fundamental in determining the schedulability of a given process system. When all processes are released simultaneously, we have the worst-case demand for the processor. The times at which all processes are released simultaneously are termed critical instants[Liu73a] (thus the first critical instant occurs at time 0). This leads to the observation that if all processes can meet their deadline in the executions starting at the critical instant, then all process deadlines will be met during the lifetime of the system. Schedulability tests for the rate-monotonic algorithm are based upon the critical instant concept. In [Liu73a] the concept is developed into a schedulability test based upon process utilisations +. The test is given by: (1) where the utilisation U i of process t i is given by: The utilisation converges on 69% for large n . Thus if a process set has utilisation of less than 69% it is guaranteed to be scheduled by the rate-monotonic algorithm. That is, all deadlines are guaranteed to be met. Whilst test (1) is sufficient, it is also not necessary. That is, the test may indicate falsely that a process system is not schedulable. For example, consider two processes with the following periods and computational requirements: For these processes, equation (1) evaluates to false, as the utilisation of two processes is 100%, greater than the allowable bound of 83%. However, when run, neither process will ever miss a deadline. Hence, the test is sufficient but not necessary. A necessary and sufficient schedulability constraint has been found [Sha88a, Leh89a]. For a set of n processes, the schedulability test is given by ++: min l T k l T k where The equations take into account all possible process phasings. 2.1. Summary In summary, schedulability tests are available for an optimal static priority scheduling scheme, the rate-monotonic scheduling algorithm, with processes limited by the following fundamental constraints: all processes are periodic; Utilisation is a measure of the ratio of required computation time to period of a process. The summation of these ratios over all processes yields the total processor utilisation. # evaluates to the smallest integer - x # evaluates to the largest integer - x # all processes have period equal to deadline; no synchronisation is permitted between processes. The first schedulability test (equation (1)) is sufficient and not necessary; the second test (equation (2)) is sufficient and necessary. One difference between the two schedulability tests lies in their computational complexities. The first test is of O (n ) in the number of processes. The second test is far more complicated: its complexity is data dependent. This is because the number of calculations required is entirely dependent on the values of the process periods. In the worst-case, the test can involve enumeration of the schedule for each process in the system, upto the period of that process. Hence, a trade-off exists between accuracy and computational complexity for these schedulability tests. The following section removes the second of the three constraints (i.e. allows processes to be sporadic thus relaxing the first of the above constraints. The third constraint is beyond the scope of this paper but as was noted earlier, Sha et al have considered this [Sha88a, Sha90a, Sha87a]. 3. DEADLINE MONOTONIC SCHEDULING We begin by observing that the processes we wish to schedule are characterised by the following relationship: computation time - deadline - period Leung et al[Leu82a] have defined a priority assignment scheme that caters for processes with the above relationship. This is termed inverse-deadline or deadline monotonic priority assignment. Deadline monotonic priority ordering is similar in concept to rate monotonic priority ordering. Priorities assigned to processes are inversely proportional to the length of the deadline [Leu82a]. Thus, the process with the shortest deadline is assigned the highest priority, the longest deadline process the lowest priority. This priority ordering defaults to a rate monotonic ordering when period =deadline . Deadline monotonic priority assignment is an optimal static priority scheme for processes that share a critical instant. This is stated as Theorem 2.4 in [Leu82a]: "The inverse-deadline priority assignment is an optimal priority assignment for one processor." To generate a schedulability constraint for deadline monotonic scheduling the behaviour of processes released at a critical instant is fundamental: if all processes are proved to meet their deadlines during executions beginning at a critical instant these processes will always their deadlines [Liu73a, Leu82a]. Using the results of Leung et al stated above as a foundation, new schedulability tests are now developed. Initially, two processes are considered, then we generalise to allow any number of processes. 3.1. Schedulability Of Two Processes Consider two processes: t 1 and t 2 . Process t 1 has a higher priority than process t 2 and so by deadline monotonic priority assignment, Consider the following case. Case (i) : both processes are always released simultaneously. This occurs if the following holds: This is illustrated in Figure 1. time time Figure 1. has the highest priority, it claims the processor whenever it has an outstanding computational requirement. This will occur for the first C 1 units of each period T 1 . The schedulability of this system is given by: (a) check schedulability of the deadline must be sufficiently large to contain the computation demand, i.e. (b) check schedulability of all higher priority processes (i.e. t 1 ) have prior claim on the processor. Hence, in any interval # can utilise the processor in the interval # # . Therefore, t 2 can have a maximum computation time defined by: that is t 2 is schedulable if The second term of equation (3) relates to the maximum time that t 2 is prevented from executing by higher priority processes, in this case t 1 . This time is termed the interference time, I. Definition 1: I i is the interference that is encountered by t i between the release and deadline of any instance of t i . The interference is due to the execution demands of higher priority processes. The maximum interference on process t i occurs during a release of t i beginning at a critical instant (by definition of critical instant [Liu73a] ). Considering the processes in Case (i), I 2 equates to the time that t 1 executes whilst t 2 has outstanding computational requirement. Thus I 2 is equal to one computation of t 1 . That is: I In Case (i) I 1 is zero, as t 1 is the highest priority process: I The schedulability for Case (i) can be restated: "i I i where I I We now consider cases where the periods of the two processes are not equal. Case is released many times before the second release of t 1 time time Figure 2. In Figure 2, the maximum interference I 2 is equal to one computation time of t 1 . The schedulability equations (4) will hold for this case. Case is released many times before the second release of t 2 Consider Figure 3. time time Figure 3. Clearly, t 2 is prevented from running by releases of process t 1 . The number of releases of t 1 within the interval # # is given by: Therefore, schedulability is expressed by: "i I i where I I In equation (5), the value of I 2 above could be larger than the exact maximum interference. This is because I 2 includes computation time required by t 1 for its (i release, some of which may occur after the value of I 2 is at least as great as the maximum interference, the test must hold since the it is based upon the exact maximum interference. An example using schedulability equation (5) is now given. Example Consider the following process system. The schedulability of the process system can be determined by equation (5). (a) check process t 1 I 1 Hence t 1 is schedulable. (b) check process t 2 I 2 I 2 where I substituting6 Hence t 2 is schedulable. An example run of the system is given in Figure 4+. Figure 4. The construction of I 2 is sufficient but not necessary as the following example shows. Consider the effect of increasing D 2 to 11. This should not affect the schedulability of the system. (c) recheck process t 2 I 2 I 2 where I substituting6 Hence t 2 is now unschedulable by equation (5). However, the process system is schedulable (Figure 4 above). Simulation diagrams are discussed in Appendix 1. The schedulability constraint in equation (5) is too strong due to the value of I 2 . An exact expression for I 2 is now developed. Consider Figure 5. time time Figure 5. A critical instant has occurred (iT with the interference on t 2 a maximum. We note that the interference consists of executions of t 1 that have deadlines before D 2 , and the execution of t 1 that has a release before D 2 and a deadline after D 2 . We can restate I 2 as I where b represents the interference due to complete executions of t 1 and k the incomplete executions. The number of complete executions in the interval # # is equal to the number of deadlines t 1 has in this interval. The number is given by: Hence, the interference due to complete executions is given by: The number of incomplete executions of t 1 is given by the number of releases of t 1 minus the number of deadlines of t 1 in # # . This evaluates to either 0 or 1. The number of releases is given by: Note that if a release of t 1 coincides with D 2 , then it is deemed to occur fractionally after . Hence the number of incomplete executions in # # is given by: The start of the incomplete execution is given by: Hence, the length of the interval utilised by the incomplete execution before D 2 is: The maximum time t 1 can use during the interval is given by the length of the interval. However, the interval may be longer than C 1 . Therefore the maximum interference due to incomplete executions is given by: Substituting b and k into equation (6) gives the following schedulability constraint: "i I i where I I Consider the following theorems which relate to the sufficient and necessary properties of equation (7). Theorem 1: the schedulability test given in equation (7) is sufficient for two processes. The proof is by contradiction. We assume there is a process system that passes the test but is not schedulable and show that if the system is not schedulable then it must fail the test. Consider a process system containing t 1 and t 2 . Let process t 1 pass the test. pass the test, but not be schedulable. To pass the test, the following must hold: For t 2 not to be schedulable, it must miss its deadline during an instance of the process starting at the critical instant of all processes. At this point t 2 suffers its maximum interference, I 2 , due to the higher priority process. Therefore for t 2 to miss its deadline and not be schedulable we have: This gives I 2 A clear contradiction exists between (a) and (b). Therefore, if t 2 passes the test, it is schedulable. The proof for Theorem 1 relies upon I 2 being exact (this is given by Theorem 2). Theorem 1 will still hold if I 2 is greater than the exact value. This merely represents a worse than worse-case. Therefore, by implication of Theorem 1, the schedulability test given by equation (5) is also sufficient. Theorem 2: the schedulability test is necessary if values of I i are exact. For process t 2 to pass the schedulability test requires: I - 2 where I - 2 represents the exact value of I 2 . When comparing I - 2 and I 2 we have three cases: (i) I - 2 > I 2 - this is clearly impossible as we know that I 2 is at least I - 2 from the above discussion. occurs when we have made a pessimistic calculation for I 2 . As I 2 increases, the computation time that could be guaranteed for t 2 decreases since: occurs when the calculation of I 2 is precise. The allowable computation time for t 2 is maximised (by above inequality). In summary, we have the greatest amount of time for t 2 if I 2 is exact. There- fore, the schedulability test is necessary if I 2 is exact. Therefore, the schedulability test given by equation (7) is necessary as the values for I i are exact. By implication, the schedulability test given by equation (5) is also necessary if I i is exact. However, I i values in equation (7) are exact in more instances than in equation (5): the former will declare more process systems schedulable than the latter. The following example illustrates this point. Example We return to the process system that failed equation (5) but was illustrated to meet all deadlines. The schedulability of the system can be determined by equation (7). (a) check process I 1 Hence t 1 is schedulable. (b) check process t 2 I 2 I 2 where I ## I substituting,6 Hence t 2 is schedulable. The system is schedulable by equation (7). A simulated run of the system was given in Figure previously. 3.1.1. Summary Noting the results stated in [Leu82a] that deadline monotonic priority assignment is optimal, two schedulability tests for two-process systems have been developed. The test in equation (5) is sufficient but not necessary, whilst the test in equation (7) is sufficient and necessary (and hence optimal). One difference between the tests is that the former is of computational complexity O # # and the latter O # # . A trade-off again exists between accuracy and computational complexity. 3.2. Schedulability Of Many Processes The schedulability test given by equations (5) and (7) are now generalised for systems with arbitrary numbers of processes. Firstly, equation (5) is expanded. Consider Figure 6. time time time Figure 6. The interference I i that is inflicted upon process t i by all higher priority processes corresponds to the computation demands by those processes in the interval of time from the critical instant to the first deadline of t i . The interference on t i by t j can be given by: This may include part of an execution of t j that occurs after D i . The total interference on t i can be expressed by: I Therefore to feasibly schedule all processes: "i I i where I Equation (8), like equation (5), is sufficient but not necessary. This is illustrated by the following example. Example Consider the following process system. The schedulability of the process system can be determined by equation (8). (a) check process t 1 I 1 Hence t 1 is schedulable. (b) check process t 2 I 2 I 2 where I substituting2 Hence t 2 is schedulable. (c) check process t 3 I 3 I 3 where I I substituting4 Hence t 3 is schedulable. Consider the effect of increasing D 3 to 11. This should not affect the schedulability of the system. (d) recheck process t 3 where I I substituting4 Hence t 3 is unschedulable by equation (8). The process system is unschedulable by equation (8). However, when the system is run all deadlines are met (see Figure 7). Figure 7. In the above example, the process system is not schedulable by equation (8) because the values of I i are greater than exact values. Each I i can contain parts of executions that occur after D i . This is similar to the drawback of equation (5) when a two-process system was being considered. To surmount this problem we generalise equation (7) for many processes. Consider Figure 8. time time time time Figure 8. From Figure 8, it can be seen that in the general case with n processes, I i is equal to the interference of all the processes t 1 to t i-1 in the interval # # . Thus, equation (7) can be rewritten to provide a schedulability test for an n process system: "i I i where I To show that the above constraint is more accurate than equation (8) consider the following example. Example We return to the process system that failed equation (8) but was shown to meet all deadlines (see Figure 7). (a) check process t 1 Hence t 1 is schedulable. (b) check process t 2 I 2 I 2 where I I ## substituting2 Hence t 2 is schedulable. (c) check process t 3 I 3 I 3 where I I ## ## I substituting3 Hence t 3 is schedulable. An example run was given in Figure 7. The expression for I i in equation (9) is not exact. This is because the interference on t i due to incomplete executions of t 1 to t i-1 given by (9) is greater than or equal to the exact interference. Consider the interference on t i by incomplete executions of t 1 and t i-1 Figure 8). Within I i , allowance is made for t 1 using all of # # , and for t i-1 using all of # # . Since only one of these processes can execute at a time, I i is greater than a precise value for the interference. Consider the following theorems. Theorem 3: the schedulability test given by equation (9) is sufficient. The proof follows from Theorem (1). Theorem 4: the schedulability test given by equation (9) is necessary if values of I i are exact. The proof follows from Theorem (2). Theorems (3) and (4) show that both equations (9) and (8) (by implication) are sufficient and not necessary. When no executions of higher priority processes overlap the deadline of t i then I i will be exact with both tests (8) and being necessary. Indeed, if the I i values in both equations (8) and are exact, the two equations are equivalent. However, when executions do overlap the deadline of t i test (9) will pass more process systems than test (8) as it contains a more precise measurement of I i . To obtain an exact value for I i under all cases requires the exact interleaving of all higher priority processes to be considered upto the deadline D i . This could involve the enumeration of the schedule upto D i with obvious computational expense. The following section outlines an alternative strategy for improving the schedulability constraint. 3.3. Unschedulability of Many Processes The previous section developed a sufficient and not necessary test for the schedulability of a process system. We note that whilst this test identifies some of the schedulable process systems, a sufficient and not necessary unschedulability test will identify some of the unschedulable systems. This approach is illustrated by Figure 9. Schedulable Systems Unschedulable Systems Systems Found By Sufficient and Not Necessary Schedulability Test Systems Found By Sufficient and Not Necessary Unschedulability Test Exact Division Given By Sufficient and Necessary Schedulability or Unschedulability Test Domain of Process Systems Figure 9. A sufficient and not necessary unschedulability test identifies some unschedulable process systems in the same manner as the test in the previous sub-section identifies schedulable systems. The combination of the two tests enables the identification of many schedulable and unschedulable process systems without resorting to a computationally expensive sufficient and necessary test. A sufficient and not necessary unschedulability test is now presented. Consider the interference of higher priority processes upon t i . This is at a minimum when any incomplete executions of higher priority processes occur as late as possible. This maximises the time utilised by higher priority processes after D i and minimises the time utilised before D i . Theorem 5 : I i is at a minimum when incomplete executions of higher priority processes perform their execution as late as possible. time time Figure 10. Consider Figure 10. The execution of t j in # # is decreased by moving the execution towards the deadline of t j (which is after movement decreases I i . I i will be at a minimum when the execution has been moved as close as possible to the deadline of t j . Consider the schedulability of t i . When I i is a minimum, we have the best possible scenario for scheduling t i . If t i cannot be scheduled with I i a minimum, it cannot be scheduled with an exact I i since this value is as least as large as the minimum value. Therefore, to show the unschedulability of a process system, it is sufficient to show the unschedulability of the system with minimum values of I i . An unschedulability test is now developed using minimum interference. In Figure 10, the interference on t i is the sum of complete executions of higher priority processes, and the parts of incomplete executions that must occur before b be the complete executions and k the incomplete executions. The total interference is stated as: I Complete executions occur in the interval # # and are given by: The incomplete executions number either 0 or 1 for each of the processes with a higher priority than t i . Hence, the interference due to incomplete executions can be stated as Substituting into equation (10) we generate an unschedulability test: "i I i where I We note that only one process need pass the unschedulability test for the process system to be unschedulable. The converse of Theorem 3 proves equation (11) to be a sufficient condition for unschedulability. This follows from the observation that since I i is a minimum (by Theorem 5), then by Theorem 5 if a process system cannot be scheduled with I i less than an exact value, the process system cannot be scheduled with exact values of I i . By Theorem 4, we note that equation (11) is a not necessary condition for unschedulability since the values used for I i are less than or equal to the exact value for I i . Equations and (11) can be used together. Consider a process system that fails equation (9). Since this is test is not necessary it does not prove the process system unschedulable. The same process system can be submitted to equation (10). If the system equation (10) we have determined the unschedulability of the process system. However, if the process system failed both schedulability and unschedulability tests we note that it could still be schedulable. We illustrate the use combined use of equations (9) and (11) with the following example. Example Consider the following process system. We can show the unschedulability of the system by using equation (11). (a) check process I 1 Hence t 1 fails the test and is therefore not unschedulable. (b) check process I 2 I 2 where I I substituting3 Hence t 2 fails the test and is therefore not unschedulable. (c) check process I 3 I 3 where I I substituting7 Therefore t 3 passes the unschedulability test. The process system is therefore unschedulable. An example run of the system is given in Figure 11. Process t 3 misses its deadline at time 13. Figure 11. We now reduce the computation time of process t 3 to 5: By observation, we can see that t 3 now fails the unschedulability test: I 3 Since the characteristics of the first two processes are identical, the process system as a whole fails the unschedulability test. However, the system is not necessarily unschedulable. Now we try to prove the process system schedulable using equation (9). (a) check process t 1 I 1 Hence t 1 is schedulable. (b) check process t 2 where I I substituting3 Hence t 2 is schedulable. (c) check process t 3 I 3 I 3 where I I substituting5 Hence t 2 is not schedulable by equation (9). In the above, we have shown the process system failing both the schedulability and unschedulability tests. Since both tests are sufficient and not necessary we have not decisively proved the process system schedulable or unschedulable. The above example illustrated the combined use of unschedulability and schedulability tests. The first part of the example utilised the unschedulability test to prove the test unschedulable. Then, by decreasing the computation time of t 3 , the system fails the unschedulability test. However, after application of equation (9), the system was shown to fail the schedulability test also. Indeed, by examining the example we can see that when C 3 lies in # # the system can be proved schedulable. When C 3 lies in # # the system can be proved unschedulable. When C 3 lies in # # we can not prove the system schedulable nor unschedulable. This requires a more powerful schedulability test. Such a test is presented in the next sub-section. 3.4. Exact Schedulability of Many Processes The schedulability and unschedulability constraints for systems containing many processes, given by equations (9) and (11) respectively, are sufficient and not necessary in the general case. To form a sufficient and necessary schedulability test requires exact values for I i (by Theorems 2 and 4). To achieve this, the schedule has to be evaluated so that the exact interleaving of higher priority process executions is known. This is costly if the entire interval between the critical instant and the deadline of process t i is evaluated as this would require the solution of D i equations. The number of equations can be reduced by observing that if t i meets its deadline at lies in # # , we need not evaluate the equations in # # . Further reductions in the number of equations requiring solution can be made by considering the behaviour of the processes in the interval # # . Consider the interaction of processes t 1 to t i-1 on process t i in the interval # # . For process t i to meet its deadline at D i we require the following condition to be met: I i We wish to consider only the points in D i upto and including t - i . Therefore, we need to refine the definition of interference on t i so that we can reason about the interval # rather than the single point in time D i . Definition 2: I i t is the interference that is encountered by t i between the release of t i and time t , where t lies in the interval # # . This is equal to the quantity of work that is created by releases of higher priority processes in the interval between the release of t i and time t . At t - i the outstanding work due to higher priority processes must be 0 since t i can only execute if all higher priority processes have completed. Hence, the point in time at which t i actually meets its deadline is given by: I i Therefore, we can state the following condition for the schedulability of I i where I i We note that the definition of I i t includes parts of executions that may occur after t . However, since the outstanding workload of all processes is 0 at t - i then when the expression I i t is exact. The above equations require a maximum of D i calculations to be made to determine the schedulability of t i . For an n process system the maximum number of equations that need to be evaluated is: The number of equations that need to be evaluated can be reduced. This is achieved by limiting the points in # # that are considered as possible solutions for t - i . Consider the times within # # that t i could possibly meet its deadline. We note that I i t is monotonically increasing within the time interval # # . The points in time that the interference increases occur when there is a release of a higher priority process. This is illustrated by Figure 12. (D 4 ) I 4 Release Release Release t 3 Release Release Figure 12. In Figure 12, there are three processes with higher priority than t 4 . We see that as the higher priority processes are released, I 4 increases monotonically with respect to t . The graph is stepped with plateaus representing intervals of time in which no higher priority processes are released. It is obvious that only one equation need be evaluated for each plateau as the interference does not change. To maximise the time available for the execution of t i we choose to evaluate at the right-most point on the plateau. Therefore, one possible reduction in the number of equations to evaluate schedulability occurs by testing t i at all points in # # that correspond to a higher priority process release. Since as soon as one equation identifies the process system as schedulable we need test no further equations. Thus, the effect is to evaluate equations in # # . The number of equations has been reduced in most cases. We note that no reduction will occur if for each point in time in # # a higher priority process is released with t i meeting its deadline at D i . The number of equations is reduced further by considering the computation times of the processes. Consider Figure 13. (D 4 ) Time 0: Release t 1 , t 2 , t 3 and t 4 . Time 4: Release t 1 Time 5: Release t 2 Time Time 8: Release t 1 Figure 13. In Figure 13 the total computation requirement of the system (C s ) is plotted against time. At the first point in time when the outstanding computation is equal to the time elapsed, we have found t - 4 (by equation (12)). In the above diagram this point in time coincides with the deadline of t 4 . Considering Figure 13, there is no point in testing the schedulability of t i in the interval # # . Also, since time 0 corresponds with a critical instant (a simultaneous release of all processes) the first point in time that t i could possibly complete is: This gives a schedulability constraint of: I i Since the value of t 1 assumes that only one release of each process occurs in # # , the constraint will fail if there have been any releases of higher priority processes within the interval # # . The exact amount of work created by higher priority processes in this interval is given by: I i The next point in time at which t i may complete execution is: This gives a schedulability constraint of: I i Again, the constraint will fail if releases have occurred in the interval # # . Thus, we can build a series of equations to express the schedulability of t i . I i I i I i I i tk If any of the equations hold, t i is schedulable. The series of equations above is encapsulated by the following algorithm: Algorithm foreach t i do while # do if I i else endif exit /* t i is unschedulable * / endif endwhile endfor The algorithm terminates as the following relation always holds. When t i is greater than D i the algorithm terminates since t i is unschedulable. Thus we have a maximum number of steps of D i . This is a worst-case measure. The number of equations has been reduced from the method utilising plateaus in Figure 11. This is because we consider only the points in time where it is possible for t i to complete, rather than points in time that correspond to higher priority process releases. An example use of the above algorithm is now given: Example We return to the process system which could not be proved schedulable nor unschedulable: proved schedulable in the previous example so we confine attention to t 3 . We use the successive equations to show unschedulability. I 3 where where I 3 substituting14 The process is unschedulable at time 12, so we proceed to the next equation. I 3 where Since we now have t 1 > D 3 we terminate with t 3 unschedulable. We reduce the computation time of t 3 to 3: We use the successive equations to show t 3 schedulable. I 3 where where I 3 substituting7 Hence t 3 is schedulable, meeting its deadline at time 10. An example run of the system is seen in Figure 14. Figure 14. The successive equations (13) have shown the process system to be schedulable. The solution of a single equation was required. 3.5. Summary This section has introduced a number of schedulability and unschedulability tests for the deadline monotonic algorithm: # a O (n ) schedulability test that is sufficient and not necessary; # a O (n 2 ) schedulability test that is sufficient and not necessary; # a O (n 2 ) unschedulability test that is sufficient and not necessary; # a sufficient and necessary schedulability test that has data-dependent complexity. The first test provides the coarsest level. The second and third tests combine to provide a finer grain measure of process systems that are definitely schedulable or definitely unschedulable. The sufficient and necessary test is able to differentiate schedulable and unschedulable systems to provide the finest level of test. One constraint on the process systems is that they must have a critical instant. This is ensured as all processes have an initial release at time 0. 4. SCHEDULING SPORADIC PROCESSES Non-periodic processes are those whose releases are not periodic in nature. Such processes can be subdivided into two categories [Aud90a]: aperiodic and sporadic. The difference between these categories lies in the nature of their release frequencies. Aperiodic processes are those whose release frequency is unbounded. In the extreme, this could lead to an arbitrarily large number of simultaneously active processes. Sporadic processes are those that have a maximum frequency such that only one instance of a particular sporadic process can be active at a time. When a static scheduling algorithm is employed, it is difficult to introduce non-periodic process executions into the schedule: it is not known before the system is run when non-periodic processes will be released. More difficulties arise when attempting to guarantee the deadlines of those processes. It is clearly impossible to guarantee the deadlines of aperiodic processes as there could be an arbitrarily large number of them active at any time. Sporadic processes deadlines can be guaranteed since it is possible, by means of the maximum release frequency, to define the maximum workload they place upon the system. One approach is to use static periodic polling processes to provide sporadics with executions time. This approach is reviewed in section 4.1. Section 4.2 illustrates how to utilise the properties of the deadline monotonic scheduling algorithm to guarantee the deadlines of sporadic processes without resorting to the introduction of polling processes. 4.1. Sporadic Processes: the Polling Approach To allow sporadic processes to execute within the confines of a static schedule (such as that generated by the rate-monotonic algorithm) computation time must be reserved within that schedule. An intuitive solution is to set up a periodic process which polls for sporadic processes [Leh87a]. Strict polling reduces the bandwidth of processing as processing time that is embodied in an execution of the polling process is wasted if no sporadic process is active when the polling process becomes occurring after the polling process's computation time in one period has been exhausted or just passed have to wait until the next period for service. A number of bandwidth preserving algorithms have been proposed for use with the rate-monotonic scheduling algorithm. One such algorithm is the deferrable server [Leh87a, Sha89b, Sha89a]. The server is a periodic process that is allotted a number of units of computation time per period. These units can be used by any sporadic process with outstanding computational requirements. When the server is run with no outstanding sporadic process requests, the server does not execute but defers its assigned computation time. The server's time is preserved at its initial priority. When a sporadic request does occur, the server has maintained its priority and can thus run and serve the sporadic processes until its allotted computation time within the server period has been exhausted. The computation time for the server is replenished at the start of its period. Problems arise when sporadic processes require deadlines to be guaranteed. It is difficult to accommodate these with a deferrable server due to the rigidly defined points in time at which the server computation time is replenished. The sporadic server [Sha89a] provides a solution to this problem. The replenishment times are related to when the sporadic uses computation time rather than merely at the period of the server process. The sporadic server is used by Sha et al [Sha89a] in conjunction with the rate-monotonic scheduling algorithm to guarantee sporadic process deadlines. Since the rate-monotonic algorithm is used, a method is required to map sporadic processes with timing characteristics given by computation time - deadline - period onto periodic server processes that have timing characteristics given by computation time - deadline = period The method adopted in [Sha89a] lets the computation time, period and deadline of the server be equal to the computation time, minimum inter-arrival time and deadline of the sporadic process. The rate-monotonic scheduling algorithm is then used to test the schedulability of the process system, with runtime priorities being assigned in a deadline monotonic manner. The next section details a simpler approach to guaranteeing sporadic deadlines based upon the deadline monotonic scheduling algorithm. 4.2. Sporadic Processes: the Deadline Monotonic Scheduling Approach Consider the timing characteristics of a sporadic process. The demand for computation time is illustrated in Figure 15. Process Released Deadline Released Deadline Released Figure 15. The minimum time difference between successive releases of the sporadic process is the minimum inter-arrival time m . This occurs between the first two releases of the sporadic. At this point, the sporadic is behaving exactly like a periodic process with period the sporadic is being released at its maximum frequency and so is imposing its maximum workload. When the releases do not occur at the maximum rate (between the second and third releases in Figure 15) the sporadic behaves like a periodic process that is intermittently activated and then laid dormant. The workload imposed by the sporadic is at a maximum when the process is released, but falls when the next release occurs after greater than m time units have elapsed. In the worst-case the sporadic process behaves exactly like a periodic process with period m and deadline D (D - m ). The characteristic of this behaviour is that a maximum of one release of the process can occur in any interval # # where release time t is at least units after the previous release of the process. This implies that to guarantee the deadline of the sporadic process the computation time must be available within the interval # # noting that the deadline will be at least m after the previous deadline of the sporadic. This is exactly the guarantee given by the deadline-monotonic schedulability tests in section 3. For schedulability purposes only, we can describe the sporadic process as a periodic process whose period is equal to m . However, we note that since the process is sporadic, the actual release times of the process wil not be periodic, merely separated by at least m time units. For the schedulability tests given in section 3 to be effective for this process system, all processes, both periodic and sporadic, have to be released simultaneously. We can assume that all the processes are released simultaneously at time 0: a critical instant. This forms the worst-case workload on the processor. If the deadline of the sporadic can be guaranteed for the release at a critical instant then all subsequent deadlines are guaranteed. An example is now given. Example Consider the following process system. are periodic, whilst t 2 and t 4 are sporadic with minimum inter-arrival times given by T 2 and T 4 respectively. We check the schedulability of the system using the equations given in section 3. The simplest test (equation (8)) is used. (a) check process t 1 I 1 Hence t 1 is schedulable. (b) check process t 2 I 2 I 2 where I substituting2 Hence t 2 is schedulable. (c) check process t 3 I 3 I 3 where I substituting2 Hence t 3 is schedulable. (d) check process t 4 I 4 I 3 where I I substituting2 Hence t 4 is schedulable. The process system is schedulable. An example run is given in Figure 16. Figure 16. In the example run (Figure 16) all deadlines are met. Each of the sporadic processes are released at time 0. This forms a critical instant and thus the worst-possible scenario for scheduling the process system. A combination of many periodic and many sporadic processes was shown to be schedulable under this scheme without the need for server processes which are required for scheduling sporadic processes with the rate-monotonic scheduling algorithm (see section 4.1). 4.3. Summary The proposed method for guaranteeing the deadlines of sporadic processes using sporadic servers within the rate-monotonic scheduling framework has two main drawbacks. Firstly, one extra periodic server process is required for each sporadic process. Secondly, an extra run-time overhead is created as the kernel is required to keep track of the exact amount of time the server has left within any period. The deadline-monotonic approach circumvents these problems since no extra processes are required: the sporadic processes can be dealt with adequately within the existing periodic framework. 5. CONCLUSIONS The fundamental constraints of the rate-monotonic scheduling algorithm have been weakened to permit processes that have deadlines less than period to be scheduled. The resulting scheduling mechanism is the deadline-monotonic algorithm. Schedulability tests have been presented for the deadline-monotonic algorithm. Initially a simple sufficient and not necessary schedulability test was introduced. This required a single equation per process to determine schedulability. However, to achieve such simplicity meant the test was overly pessimistic. The simplifications made to produce a single equation test were then partially removed. This produced a sufficient and not necessary schedulability test which passed more process systems than the simple test. The complexity of the second test was O (n 2 ) compared with O (n ) for the simple test. Again, the test was pessimistic. To complement the second schedulability test, a similar unschedulability test was developed. The combination of sufficient and not necessary schedulability and unschedulability tests was shown to be useful for identifying some unschedulable systems. However, it was still possible for a process system to fail both the schedulability and unschedulability tests. This problem was resolved with the development of a sufficient and necessary schedulability test. This was the most complex of all the tests having a complexity related to the periods and computation times of the processes in the system. The complexity was reduced substantially when the number of equations required to determine the schedulability of a process were minimised. The problem of guaranteeing the deadlines of sporadic processes was then discussed. Noting that schedulability tests proposed for sporadic processes and the rate-monotonic algorithm require the introduction of special server processes, we then proposed a simple method to guarantee the deadlines of sporadic processes within the confines of the deadline-monotonic algorithm. The simplicity of the method is due to sporadic processes being treated exactly as periodic processes for the purpose of determining the schedulability. Using this scheme, any mixture of periodic and sporadic deadlines can be scheduled subject to the process system passing the deadline-monotonic schedulability constraint. A number of issues raised by the work outlined in this paper require further consideration. These include the effect of allowing processes to synchronise and vary their timing characteristics. Another related issue is the effect of deadline-monotonic scheduling upon system utilisation. These issues remain for further investigation. ACKNOWLEDGEMENTS The author thanks Mike Richardson, Alan Burns and Andy Wellings for their valuable comments and diatribes. --R "Misconceptions About Real-Time Computing: A Serious Problem for Next Generati on Systems" "Scheduling Real-Time Systems" "Scheduling Algorithms for Multiprogramming in a Hard Real-Time Environment" "The Rate-Monotonic Scheduling Algorithm: Exact Characterization and Average Case Behaviour" "Aperiodic Task Scheduling for Hard Real-Time Systems" "Real-Time Scheduling Theory and Ada" "Real-Time Scheduling Theory and Ada" "Priority Inheritance Protocols: An Approach to Real-Time Synchronisation" "On the Complexity of Fixed-Priority Scheduling of Periodic, Real-Time Tasks" "Enhanced Aperiodic Responsiveness in Hard Real-Time Environments" "An Analytical Approach to Real-Time Software Engineering" --TR

deadline monotonic scheduling;priority inheritance protocol;rate monotonic scheduling;feasibility decision algorithm;periodic task

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The Synchronous Approach to Designing Reactive Systems.

Synchronous programming is available through several formally defined languages having very different characteristics: Esterel is imperative, while Lustre and Signal are declarative in style; Statecharts and Argos are graphical languages that allow one to program by constructing hierarchical automata. Our motivation for taking the synchronous design paradigm further, integrating imperative, declarative (or dataflow), and graphical programming styles, is that real systems typically have components that match each of these profiles. This paper motivates our interest in the mixed language programming of embedded software around a number of examples, and sketches the semantical foundation of the Synchronie toolset which ensures a coherent computational model. This toolset supports a design trajectory that incorporates rapid prototyping and systematic testing for early design validation, an object oriented development methodology for long term software management, and formal verification at the level of automatically generated object code.

Introduction Reactive computer systems continuously respond to external stimuli generated by their en- vironments. They are critical components of our technology dominated lives, be they in control systems such as ABS for cars, fly-by-wire in aircraft, railway signalling, power gen- eration, shopfloor automation, or in such mundane things as washing machines and video recorders. Mastering the design of these systems, and reducing the time needed to bring them to market, becomes of utmost economic importance in times of increasing market dynamics. This paper advances the paradigm of synchronous programming as a means to match these goals. Several programming languages that originally emerged as engineering notations are now defined in the IEC Standard 1131-3 [18]. These languages have been designed specifically for embedded software applications, mainly in the process control industry but with increasing influence in other sectors. Unfortunately the IEC 1131 languages have not been designed with the benefit of formal semantics. Yet, to advance the state of the practice of embedded software design it is important to provide tools that support high-level specification and rapid prototyping, integrate testing and formal verification to achieve early design validation, and encourage modular software development-to ease review, maintenance, and certification. For this, languages having precise mathematical semantics are required. Synchronous programming languages like those discussed in this paper have the potential to introduce such rigour embedded software design. These languages have very distinctive characteristics: ESTEREL [4] is imperative in style, while LUSTRE [13] and SIGNAL are declarative; STATECHARTS [15] and ARGOS [26] are graphical notations that enable one to program directly by constructing hierarchical automata. These languages share a common communication metaphor, that of synchronously broadcast signals. Sections 2 and 3 introduce the key ideas behind the synchronous approach to embedded soft- ware, and outline some of the programming constructs available in LUSTRE, ESTEREL, and ARGOS through a few simple examples. A longer introduction to these languages can be found in [12]. The specific profiles of these languages reflects the fact that they have been developed in response to problems emerging in different application areas. LUSTRE and SIGNAL derive from requirements of industries mainly aware of electrical and electronics engineering methods, who wanted to manage the increasing complexity of their applications, and gain greater flexibility in their design, by introducing software. These languages have therefore been designed for the discrete handling of continuous phenomena. They invoke metaphors commonly used by electrical engineers in control theory, and are thus most suited to signal-based applications such as in navigation or digital signal processing where often a sampling of different related frequencies can be found. In contrast, ESTEREL and graphical languages like ARGOS are better suited for scheduling complex operating modes, handling intricate patterns of events, and describing interrupt-driven behaviours. However, the difference in their best-suited application profiles is only a qualitative assessment since these languages have broadly similar expressive power. Our motivation for taking the synchronous design paradigm further, with the wholesale integration of imperative, declarative (or dataflow), and graphical programming styles, is that real systems typically have components that match each of these profiles. As argued by Gajski [11], the construction of embedded systems requires a combination of state based and dataflow models which support hierarchical structuring of behaviour, concurrency, and exception handling. We believe a semantical combination of the languages described here will definitely satisfy these requirements. That is the unique feature of the SYNCHRONIE workbench which is under active development in the Embedded Software Design group at GMD. The success of this venture rests on the underlying computational model sketched in Section 4, and with greater rigour in Section 5, which provides a coherent mathematical framework, and yields compact, verifiable code. The main functions of the SYNCHRONIE workbench are described in Section 6. Synchronous programming languages are already being evaluated in some industries, particularly aviation and power generation where the problems are real-time constrained and safety critical. The commercial interest in synchronous languages lies not only in the style, but in their seamless integration with existing software development practices through the programming environments with which they are provided. A short overview of the marketed environments that support each language, and some industrial applications, can be found in [2, 19]. In addition to project management facilities, and editing facilities that mix text and graphics, the commercial environments provide advanced features for design vali- dation, and back-end compilation to various languages like C or VHDL. The SYNCHRONIE workbench also provides such features, but offers system designers in addition much greater freedom in the choice of programming language, with the unique option to freely mix the various synchronous programming modes. For the most part synchronous languages provide primitive datatypes and operations on them only: the emphasis is on gaining intellectual control over the program's often intricate control logic, rather than on data processing issues. Where compound datatypes need to be used, the synchronous language is 'hosted' in a common language such as C, Fortran, or Ada. Synchronous languages achieve good separation between concerns of data and control logic; their novel fusion in SYNCHRONIE with object oriented construction techniques [7, 8] offers the right kind of design encapsulation and abstraction mechanisms to achieve much needed transparency in the software development lifecycle. 2. Reactivity and Synchrony In contrast to an interactive system (say a text editing program), a reactive system is fully responsible for the synchronisation with its environment. A system is reactive when it is fast enough to respond to every input event, and its reaction latency is short enough that the environment is still receptive to its responses. Most control systems, and systems for digital signal processing as used in industry or in telecommunications, are reactive according to this characterisation. Common features of reactive systems are: concurrency: they typically consist of several concurrent components that cooperate to realize the intended overall behaviour. real-time: they are supposed to meet strict constraints with regard to timing, such as response time or availability. determinism: a system's reaction is uniquely determined by the kind and timing of external stimuli. heterogeneity: they often consist of components implemented in quite different technologies like software, hardware, or on distributed architectures. reliability: requirements include functional correctness as well as temporal correctness of behaviour; also robustness, and fault tolerance. The fundamental idea of the synchronous approach is simple: reactive systems are idealised by assuming that stimulation and reaction are simultaneous, or that reaction takes zero time-meaning no observable amount of time. A system is stimulated by events from the environment, but responds instantaneously. Physical time does not play any special role: time will be considered as a sequence of a particular kind of external events. From this point of view the statement "the train must stop within 10 seconds," is not essentially different to the statement "the train must stop within 100 metres." In both cases something is said about the occurrence of events [12]: "The event stop must precede the 10th (or 100th) next occurrence of the event second (or metre)." Events are manifest by signals which are broadcast throughout the system instantaneously. A system reacts by emitting (that is, broadcasting) signals as well, so the statement above could be modified to: "The signal train stops must be emitted before the signal second has been emitted Abstracting physical time offers a number of advantages: the granularity of time may be changed without affecting the sequence of events, and system components can be composed and decomposed into subcomponents without changing the observable behaviour. This is particularly beneficial for proving system properties. In practical terms, the synchrony hypothesis states that a system reacts fast enough to record all external events in the proper order. This property is realistic only insofar as it can be checked: it corresponds to the hardware point of view that the time needed by an operation is of no importance as long as it does not exceed the duration of a clock cycle, or, vice versa, that the clock cycle is determined by the operation which consumes most time. The synchronous approach therefore advances a two-stage design for reactive systems: physical time is first abstracted to focus on the functions to be maintained, the validity of this abstraction being actually verified when these systems are implemented. 3. Synchronous Languages In this section we sketch some typical synchronous programming styles, focusing in the subsections that follow on ESTEREL, ARGOS, and LUSTRE. A larger example is described in Section 4.1 that uses these languages together to solve a programming problem in a natural way. Here we wish to give the reader a feeling for synchronous programming in general, so our examples are necessarily rather simple. 3.1. Imperative Programming in ESTEREL Suppose the following informal specification has been given: "If a second mouse click succeeds a first one within 5 milliseconds there is a double click, otherwise there is a single click." Assuming that time units are specified by the signal tick, the required behaviour is captured by the ESTEREL program (not the shortest) in Figure 1. At first the program waits until a click signal is produced by the environment. Then two programs are executed in parallel. The upper subprogram broadcasts the single signal after five ticks provided that loop trap done in await click; await 5 tick; present double else emit single; exit done await click; emit double; exit done loop Figure 1. An ESTEREL mouse controller double is not emitted simultaneously; it then raises the exception (the trap signal) done, and terminates. The conditional programming construct present . end should be interpreted as 'if the double signal is present do nothing, otherwise emit single (and exit)'. The other subprogram emits the double signal if click occurs a second time, and exits. Trap signals are exceptions which, when raised, abort all programs within the scope of their declaration. The net effect of raising done in either of the parallel branches in the mouse program is that the body of the loop is terminated; ESTEREL's semantics cause the loop to restart immediately, so the program returns to the await click at the top of the loop. Note the priority given to double over single if the second click should happen at the fifth tick of the clock. In contrast to the trap statement, the construct do halt watching click (which is what await click actually abbreviates) involves a second, stronger kind of preemption mechanism. The halt statement is the only one in ESTEREL to consume time; in fact halt starts but never terminates. However, the body of the watching construct will be preempted whenever the watchdog condition becomes true (that is, whenever the signal click is present in this example). All language constructs of ESTEREL are instantaneous apart from the halt statement, or derived constructs like await. As we saw, the loop restarts immediately it terminates; sequential composition is likewise reckoned to take zero time, as are the tests in conditional statements. Thus, if a second click does coincide with the fifth tick, the present test in the first subprogram in the example above will be executed in the same logical instant as the click which aborts the await in the second subprogram. Of course, no assumptions are made here about behaviour with respect to physical time which is represented explicitly by the tick signal supplied be the environment. 3.2. Graphical Programming in ARGOS To illustrate a graphical notation the program of the previous section is now coded in ARGOS (see Figure 2). States are represented by rounded boxes. Automata are hierarchic in that states can contain subautomata (for example in state two there are two subautomata). Automata can run in parallel, which is indicated by a dashed line. Finally, the scope of signals can be restricted-this is indicated by a square cornered box, instead of a rounded one, mouse# done two four three one click/ double,done click timeout five timeout.-double/ single,done# done Figure 2. ARGOS program for the mouse controller with a list of signals glued to it. A default arrow indicates the initial state of each subau- tomaton (e.g., those labelled one, three, and Count5). When started the mouse automaton enters the initial state (labelled one here). Occurrence of the click signal causes a jump from there to state two, and this initialises both of the parallel subautomata. A label on an arrow that is of the form # indicates that the transition should take place if signal # arrives, and that this will cause the signal # to be emitted simul- taneously. Generally the guard ( # ) on a transition can be a list which specifies the presence or absence of a number of signals, and the output action ( # ) is a list of signal names (omit- ted, if the list is empty). So the (subsequent) transition from state Count5 to state five will fire if the timeout signal is present and the double signal is not present; this causes both single and done to be emitted. We have to add that the refinement of state Count5 contains the subautomaton displayed in Figure 3. At the fifth tick the timeout signal is emitted and the subautomaton comes to a tick tick tick tick/ timeout tick Figure 3. ARGOS 5 tick counter (i.e., await 5 tick; emit timeout) halt (tick is implicit in the guard on every transition in ARGOS, but is mentioned explicitly here if the guard would otherwise be empty). Meanwhile, if a second click occurs the transition from state three to state four takes place causing the simultaneous emission of double and done. However, if the second click occurs at the fifth tick the transition to state five is not possible because double occurs negatively in the guard of the transition. Hence, exactly one of single or double will be emitted from state two, and in either case this will be accompanied by the (local) signal done. This signal causes state two to be abandoned when it is emitted because it appears in the guard on the transition out of that EI0BKJML > N I0O 90> 9;P JQF 88 OF state-whatever individual states the subautomata are in, when this signal occurs they will be abandoned. Such control flow is similar to the preemption mechanisms of ESTEREL. The program therefore returns to state one, and waits for the next click. 3.3. Declarative Programming in LUSTRE l mn l l l prq s s ev e s prq s Figure 4. A first order digital filter To illustrate the declarative synchronous programming style we shall implement a recursive digital filter in LUSTRE. A first order digital filter may be specified by a signal flow graph [32] such as that in Figure 4. Quantities on incoming edges at the nodes of this graph are summed, and their result is broadcast along the outgoing edges. The labels pq?s , ev , e s and s on the arcs denote a delay by a shift register, and multiplications by the respective constants. In linear form we would have the equation uyx/z ev wx{ e s wx q?s uyx q?s where |T}~ denotes the time index. The boundary condition is z ev v . This equation with its boundary condition translates to the LUSTRE node: node FILTER let (one has to instantiate the constants to specific real values of course). The equation is evaluated on every program cycle which is marked in LUSTRE by an implicit clock tick, rather than the explicit tick seen earlier. Once a LUSTRE program is started it runs forever without terminating, executing every equation once a cycle. In this case only at the first tick will the term a0 * x be evaluated; at all later times the term a0 * x b1 * pre(y) is evaluated. This gives the semantics of the followed-by operator ->. The pre operator allows access to the previous value of the expression on which it operates. LUSTRE has other operators for upsampling and downsampling. These operations are illustrated in the timing diagram below. Downsampling is by means of the when operator, and upsampling is by means of the current: - . - . - . takes the value of X only if the 'clock' signal B is true (it is undefined otherwise), and current(Z) latches the value of Z up to the next sampling signal, i.e., the next instant B becomes true (the # is explained in Section 5). This mechanism allows one to easily define the digital filter with regard to a different base frequency: node BFILTER (b : let Other declarative synchronous languages include SIGNAL and SILAGE [22, 16]. The latter is mainly used as the specification language of the CATHEDRAL [21] toolset for the synthesis of DSP chips. These languages are quite similar to LUSTRE in style but use slightly different mechanisms for upsampling and downsampling. This concludes our survey of the synchronous programming styles. Most of the features of ARGOS and LUSTRE have been mentioned, but ESTEREL is a slightly richer language than the mouse controller example illustrates since, in particular, signals can carry data which may be tested to modify the control flow. These languages share the communication principle of synchronously broadcast signals, with the scoping mechanisms shown. How- ever, the fact that signals can be tested and emitted simultaneously in the parallel branches of a synchronous program can sometimes give rise to confusion over the cause and effect of a signal's activation. Such causality cycles (they correspond to short circuits in sequential are programming errors that can be detected statically by the compilers; they do not affect the semantics of the individual languages as such, only the class of acceptable programs. The detection of causality cycles is intricate however, and falls outwith the scope of the present paper (see [25, 35], for example). 4. Integration of Synchronous Languages Since complex systems often have components which match each of the different profiles sketched above, it is natural to wish to express each in the most appropriate language rather than shoe-horning solutions from a single language. To illustrate the useful interaction between these synchronous languages we reprogrammed the production cell case study [23] that has been used for evaluating and comparing software development methodologies and tools. The successful design and verification of the controller has been discussed independently in the contexts of ESTEREL, SIGNAL and LUSTRE [23], so we do not dwell on the details of the specification in Section 4.1 below, but rather on its overall organisation. Section 4.2 outlines some aspects of the underlying computational model which makes possible the idea of integrating these languages into a single, coherent mathematical framework. The formalities are drawn out in Section 5. 4.1. Multi-modal Programming The production cell's input arrives via a feed belt that conveys metal plates to an elevating rotary table; the table lifts each plate to a position where a robot picks it up with its first arm, transferring it to a press. When the plate has been forged by the press it is removed to a deposit belt by the second arm of the robot. The circuit completes with a crane that unloads the forged plates at the end of the deposit belt. Although the circuit for one item is quite simple, the design of the cell's control software should maximise the throughput while meeting the various constraints of the devices. The cell's controller has a short initialisation phase during which the various devices are set to specific (safe) states, followed by an endlessly looping process that controls the actual production cycle. The production phase divides naturally into six components that run in parallel, one for each physical device. To represent the design at this highest level of organisation we used ARGOS as illustrated in Figure 5. The graphical nature of this lan- Production Cell done Production deposit belt crane feed belt table robot press Initialise iCtrl iData done Figure 5. The Production Cell in ARGOS guage makes it very suitable for presenting high-level design choices, and the reader can appreciate at a glance the overall structure of the control program. The graphical style eases communication between partners in the development of the software, and others involved in constructing the production cell including management, mechanical and electrical engineers designing the physical components, system and safety analysts, and so on. The initialisation phase sets both arms of the robot to a retracted state, ready to handle the first plate to arrive (for safety since they might otherwise be damaged in the press when it is switched on). This behaviour is implemented by first emitting signals Rretract1 and Rretract2 to the actuators of the robot, and then waiting for the arms to reach the desired positions. Reaching these positions is signalled by OutPress1 and OutPress2, and when this occurs a stop command is emitted to the respective arm by means of the signals Rstop1 and Rstop2. When both arms have reached the desired positions the signal done is emitted, transferring control to the unending production phase. If we were programming only in ARGOS (or STATECHARTS, for that matter), there would be little choice but to implement this initialisation logic in a program similar to that in Figure 6(a). A drawback of the graphical formalism becomes apparent. The control flow is iCtrl OutPress1/ tick/ OutPress2/ tick/ Rstop2/ done Rstop1/ done Rstop1. Rstop2/ done Rstop1 (a) ARGOS initialisation module iCtrl: input OutPress1,OutPress2; output Rretract1,Rstop1, Rretract2, Rstop2, done; emit Rretract1; await OutPress1; emit Rstop1 emit Rretract2; await OutPress2; emit Rstop2 emit done module (b) ESTEREL initialisation Figure 6. Initialisation phase of the Production Cell confused by the fact that one has to explicitly manage the synchronisation to emit done when both arms have been retracted (via the process in the lower part of the figure). Also, the reader may be forgiven for wondering whether some external process running in parallel with Initialise can interfere by emitting an Rstop1 signal, say. (There is no such process, but this cannot be inferred simply by looking at Figures 5 and 6.) At this level of detail graphical programming becomes cumbersome, and therefore error prone, and one quickly loses sight of the flow of information. Instead, we can refine the state iCtrl in Figure 5 with the ESTEREL program shown in Figure 6(b). This makes the natural control flow explicit (with the semi-colon after the concurrent initialisation of the two arms), and ESTEREL's powerful parallel operator handles the synchronisation on our behalf to ensure that done is only emitted when both branches have terminated. This program only handles pure signals (that is to say, those signifying events) and no data-even though the position of the arms are in reality provided by a potentiometer delivering a real value. In the next section we shall describe one method of handling these data (refining state iData) to complete the initialisation program. The rest of the program is implemented in LUSTRE as described by Holenderski [17], and will not be discussed here. The steady state behaviour of the press, the robot, and the other physical components is adequately expressed in LUSTRE since these are interlocked, nonterminating parallel processes. However, specifying the sequential composition of the initialisation and production phases in LUSTRE leads to an obscure program since the sequential composition operator is simply not available. To avoid this in his implementation, Holenderski programmed the sequence at the level of the C interface (to the simulator provided for the case study), but this ad hoc approach is highly error prone in general, undermines the formal definition of the synchronous language (LUSTRE, in this case), and pre-empts our performing any formal verification of the full program. 4.2. Synchronous Automata Performing the sequential composition in ARGOS is, in contrast, fully formal as long as the meaning of the combination of the synchronous languages is clear. We are aided here by their relative simplicity (e.g., when compared to Ada or C), and though they each have a very different 'look and feel', ESTEREL, LUSTRE, and ARGOS can be interpreted in the same computational model. This section introduces the main notions behind synchronous automata, the more formal presentation being deferred until Section 5. 4.2.1. Boolean Automata Boolean automata are easier to understand and are, in fact, a particular instance of synchronous automata which capture the essence of the synchronous languages. Boolean automata have two kinds of statements: s signal s is emitted if condition # is satisfied, and the control register h is true (or 'set') in the next instant if # is satisfied. A synchronous program is represented by a collection of such statements defining signals to represent transient information, or registers to represent persistent information. The operational semantics of Boolean automata is defined by two successive phases: given a val- uation # that assigns a truth value, tt or ff, to each of the registers and inputs (inputs are represented by free variables), a reaction is a solution of the system of equations this solution extends # to cover all signals, and we use this to compute the assignments # to yield the next state of the machine. A solution to the signal equations for all input patterns and (reachable) states must be proved to exist at compilation time to guarantee that the program is reactive in that it may respond to every input stimulus. In addition, this solution must be unique to guarantee that the program is deterministic. These issues are common to all synchronous programming languages. For an example, the behaviour of the ESTEREL statement await OutPress1 is defined by the Boolean automaton: The register h # captures the pause in this await construct. All registers are initially false; # will be set when the Boolean automaton is initialised-that is, when the special start signal is present (it is only present, or true, in the very first clock cycle). Thereafter, h # is set every program cycle until the signal OutPress1 occurs, at which point the automaton terminates (h becomes, and stays, false). Termination is signified by the special signal # . Note that the OutPress1 will be ignored in the very first clock cycle-readers familiar with ESTEREL will realise that an 'immediate' is needed to handle that event. The Boolean automaton captures the behaviour of the statement emit Rstop1: this automaton emits the signal and terminates immediately. The sequence await OutPress1;emit Rstop1 specifies that if the former statement terminates, control passes instantaneously to the latter. This control flow is tracked by the compiler which substitutes for the start condition ( # ) of the Boolean automaton associated with the second command in sequence, the termination condition of the first. Hence, for the upper parallel branch of Figure 6(b), we obtain: We cannot sketch the translation of all the language constructs here, but have hopefully provided some feeling for how the translation proceeds. ARGOS and Boolean LUSTRE also have natural and compact interpretations in Boolean automata. The full translation of pure ESTEREL, along with its proof of correctness with respect to the published semantics, is given in [34]. 4.2.2. Synchronous Automata Synchronous automata represent an enhancement to the model to handle data; Boolean automata only capture the pure control of synchronous languages, focusing on synchronisation issues and ignoring the data carried by signals, and actions upon them. Intuitively, this enhancement is achieved by coupling the presence of a signal with a unique datum. The earlier notion is refined thus: s # is satisfied, the signal s is emitted with the value returned by the function , and # is satisfied, the register h is set for the next instant with the result of the function # . In this framework, Boolean automata are synchronous automata where the domain of values is restricted to a single-point set denoted # . Such signals are referred to as pure signals (also pure, or control, registers) because the values are of little interest. So, the earlier Boolean equations (s # ) and assignments (h # ) have now given way to conditional equations (s guarded commands (h automata are thus pure synchronous automata. To illustrate the idea let us specify the discretisation of the position of the arms of the robot in the production cell in the initialisation phase thus: node iData (Arm1, Arm2 : real) returns let This LUSTRE node emits OutPress1 and OutPress2whenever the sensors on the arms of the robot indicate that they are a safe distance from the press. This is defined by the synchronous automaton: OutPress1 OutPress2 idata The first two statements translate the equations of the LUSTRE node; the latter represents the implicit control in declarative programs. The register h idata is initially inactive, gets set (or becomes active) when the program is started with # , and remains set thereafter. This persistence is indicated by h idata appearing in the statement guard where it refers to the active status of the register, not its value. Since it never again changes state, h idata might be thought redundant. Nevertheless, it captures the nonterminating property of declarative programs: once a declarative program gets started via retains control in that it is executed at every later instant of time. We anticipate that it may lose control as well. The idea required to integrate declarative and imperative styles is formally captured in synchronous automata by the control axiom. 4.2.3. The Control Axiom The control axiom states that no synchronous automaton can react if it is not in control- a synchronous automaton has control either if # is present, or some of the registers of the process are set. By way of a non-example, the following statement alone does not define a synchronous automaton since the presence of OutPress1 neither depends on # , nor any state (the latter dependency exists implicitly in a LUSTRE program's main loop): OutPress1 The control axiom appears to contradict the notion that a reactive system must maintain an ongoing relationship with its environment, and thus must always react. But even if this is true for a complete reactive system (the whole program), some of its subcomponents may be active for just a while; in particular, they may be preempted. In the production cell, while the ESTEREL module iCtrl terminates by itself after emitting done, the LUSTRE node iData never terminates. But, with respect to the semantics of ARGOS, the Initialise state Figure is preempted by done, which therefore aborts the LUSTRE node. This is performed (in the compiler) by redefining h idata thus: idata done h idata This guarding of register statements translates a weak preemption mechanism [3]. This is similar to ESTEREL's trap construct, and is the only preemption mechanism of ARGOS. The strong preemption mechanism is captured by guarding both kinds of statements. A longer presentation of synchronous automata and their algebra is given in [24], whence the summary of the formal semantics in the next section is drawn. 5. Semantics in a Nutshell It should be evident by now that the synchronous programming paradigm has a dual nature. On the one hand languages such as LUSTRE are descriptive in that they constrain possible behaviours; one the other hand languages like ESTEREL and ARGOS foster a constructive point of view in that an automaton is specified which prescribes the transformation of a given state into another. We argued informally above that synchronous automata are a useful intermediate representation. These structures are reasonably unsophisticated mathemat- ically, and match both programming styles well enough to serve as a kind of synchronous object code, though of course originating in the tradition of deterministic automata. While not wishing to lose the light presentational style of the previous sections, or burden the reader with undue formality, we shall try to shed some light on the formal semantics in a slightly more rigorous manner. The subsections that follow first address the descriptive, then the prescriptive aspects of synchronous programming, before merging 'dataflow' and 'controlflow' in their synthesis at the end. 5.1. The Declarative Aspect: Constraining Dataflow Behaviour is manifest in what it is possible to observe. We classify observations by attributing a name which, for the moment, will be referred to as a signal. A signal may be present with a value taken from a set # , or it may be absent. If a particular signal is observed over time, a flow of values and 'misses' (when the signal is absent) will be obtained: # . EI0BKJML > N I0O 90> 9;P JQF 88 OF A flow is characterised by a subset - of the natural numbers (the sampling rate of a valuation . If |- - is present at instant | , and takes the value -|0- . In the example above, z -D~-y-0!-y- - . We refer to a set of flows as a synchronous process, and at each instant of time require that at least one signal of a synchronous process is present. This requirement allows us to identify 'observed' (or external) time with the natural numbers. Conceptually, synchronous dataflow programs deal with the specification of such pro- cesses. For example, the LUSTRE equation specifies constraints upon the corresponding flows: u the sampling rates are the same, and at each instant p the values are the same. Taking a second example, the memoisation operator pre in Y = pre(X) introduces a delay so that the value of X observed at one sampling point is observed on Y at the next: - . pre(X) - . At the first sampling point no value of X can have been observed beforehand so we insert a 'non-value' which indicates that even if Y is present, its value is undefined. Of course, a compiler must guarantee that no program's reaction ever depends on a non-value. This phenomenon is similar to that of program variables (in other languages) which should not be used at run-time before they have been initialised. LUSTRE's `followed-by' operator, discussed in Section 3.3, allows one to initialise signals properly. The formal definition (of pre) is quite elaborate since it depends on the sampling points. Skipping the exact definition, let us use the notation - w- to refer to the th sampling point in sequence, counting from 0. Then: - , and s - , for all sampling points w- , and z Synchronous automata are introduced here as a more elementary language for specifying constraints on dataflows. The statement specifies that, whenever the pure signal - is present, the signal must be present and its value must be equal to that of w . In more formal terms: u , and for all Notation is somewhat overloaded here. To be precise, each subset - IN is in one-to-one correspondence with its characteristic function - such that z tt iff |- . Hence the Boolean operators used in control expressions (the guards after the @ -\ 89;:=<?>A@ B0CDBE/F0GH are well justified, and can be unambiguously interpreted as Boolean operators, or as operators on subsets of IN. A point to note is that although the definition above forces u to have a value at instant only if may still have values at other instants. Furthermore, we require the signal w to be present at instants |- , otherwise the constraint The idea naturally generalises to more complex statements - where - is a data expression and - is a control expression. Closely related to the pre operator is the second type of statement found in synchronous automata, namely This delays the observation on w by just one instant, but upsamples at the same time: when- ever |- , whatever was observed on in instant | will be observed on u at instant | { - . The figure s interprets the more formal definition u , and for all { -y- z-w -=|0- where z -| { -|-- . However, it is not a totally trivial exercise to prove that, provided attention is restricted to signals w and u , the synchronous automaton init z s s s is equivalent to the LUSTRE statement pre(X). This introduces the third (and final) clause in synchronous automata, viz init which is used to initialise a dataflow that would not otherwise take a value 'at z While the proof that the above automaton implements pre is a little tricky, the reader can easily verify that the semantics compute the flows in the adjacent diagram which shows how the 'last' value of w is stored until the next sampling point. Note that t , so it turns out that z IN. For the diagram we supposed - . EI0BKJML > N I0O 90> 9;P JQF 88 OF 5.2. The Imperative Aspect: Managing Control The imperative synchronous languages are inherently based on the idea of a (distributed) state. In ESTEREL, for instance, the halt statement is used to indicate local sequence control within the parallel branches of a program. The halt's behaviour is defined by - ff which abbreviates where an is introduced as a specific control register for each such halt. Once this synchronous automaton has been activated by , the register remains active forever, and the automaton never terminates. This is precisely reflected in the formal semantics of which yields -/- that z -~r- this means -z- . In this discussion we revert to Boolean automata because control is specified in terms of pure signals only. We apologise for any terminological confusion that may be troubling the reader at this point: we stated earlier that flows represent observations made of certain at- tributes, namely signals, but now speak in terms of registers. To reconcile the nomenclature we shall also relate flows to registers, but stipulate that registers are 'unobservable' (as one might expect). Registers may be active (and thus possessing a value, though it may be ' in the case of control registers), or inactive. In order to model preemption in synchronous automata the simple idea is to guard 'reac- tivation' in that the above automaton is modified to: --Once the control register is activated, it remains activated only up to the instant that the signal is present; the termination signal - is emitted at the same later instant. Note that - requires that 'control' is with - (i.e., - is active). These clauses capture precisely the semantics of ESTEREL's await statement (cf. Section 3.1) where await click is defined as do halt watching click This construct generalises to do P watching s, where P is an arbitrary program state- ment. In addition to the preemption of halts, ESTEREL's semantics require that no signal is emitted by P if the watchdog signal - is present; also, - is to be ignored by P in the instant that it commences execution [4]. The preemption effect can be achieved by guarding the emittance of signals. A typical clause in P that depends on a register - may be -=e -V-0-R-=g meaning that - is emitted if e is present in the first ( or if the control register - is active when g is present. The emission of - is preempted by the watchdog action thus: -=e -V-0-R-=gV- -Note that - is still emitted in the first instant if e is present. In contrast to this, ARGOS only allows preemption of registers. With regard to Figure 2, emmittance of the signal done should preempt all the states inside state two, including state two itself. The reader will note that preemption of signals must be avoided because otherwise the signal done, being the cause of the preemption in this case, would not then be emitted. Preemption of registers is achieved by the same guarding mechanism as before, now with the guard done. Working through the details, the reader may quickly check that h one done # h two click h two click done # h two relates to the transitions from state one to state two, and conversely. Further, building on the informal description in Section 3.2, it is easy to check that the behaviour of the subprocess in state two is captured by the synchronous automaton double click single # timeout double done # timeout double click h three click # timeout double (omitting the register clauses refining state Count5, and the timeout signal). We have not defined h four or h five here since these 'terminal' states in the respective ARGOS subautomata never become activated. Guarding all the registers of # with done therefore (the ' clauses do not change, so are not repeated), we obtain h one done # h two click h two click done h two h three click click done click # timeout double done h Count5 which is a fairly precise account of what wasdescribed informally in Section 3.2. Of course, a more structural, or compositional treatment is needed for an ARGOS compiler to synthesis # from # in general. Thus, for example, the ARGOS compiler must bind h four and h five since it is not until the context surrounding these subautomata is known that it becomes clear that these registers never become activated. A compositional semantics for ARGOS in terms of synchronous automata which allows us to make such optimisations as that above is rather easy to establish, but the same cannot be said for the compositional semantics of ESTEREL, which is a much more complex lan- guage. However, to give a flavour of the approach let us return to the discussion of sequential composition begun in Section 4.2. Assume that P and Q are translated to synchronous automata # and # respectively. We use the common trick when modelling P;Q within a parallel calculus (see, e.g., [29]) of placing # and # in parallel, but preventing # from starting until signifies that it has terminated. The start signal suspends (by the control axiom) the activation of # emits its The notation # means that # is given a fresh name # in # (the same replaces # in # ). The beauty of synchronous automata is that the ' just textual juxtaposition (concatenation). Since # is broadcast, # starts in the same instant terminates. Formally, termination (of # ) requires that all control registers become inactive at the end of the instant in which emits # . This is a defining property of synchronous automata, like the control axiom discussed in Section 4.2.3. Fortunately, these turn out to be invariants of our compositional semantics of ESTEREL. Specifically, in [34] we prove that given # is the translation to synchronous automata of the ESTEREL statement P: 1. if emits # , then in the next instant no control register # in # is active, and 2. if no control register # in # is active, then no signals are emitted from # . These results support the main theorem in [34] which states that our synchronous automata semantics coincides with the official behavioural semantics of (pure) ESTEREL [4]. 5.3. Merging Control and Dataflow The basic question is: what is control in dataflow? It seems that calling a dataflow program magically transfers 'control' to the program, just as killing the resulting (endless) process withdraws control. To get a grasp of the issue, recall that "a system is reactive when it is fast enough to respond to every input event, and its reaction latency is short enough that the environment is still receptive to its responses." Implicit in this point of view is the notion that without an input event, no 'response' is given-that is: if some signal is present then some input signal must be present. Hence, it should be possible to control a synchronous process by guarding the input. Let X and Y be input signals to the LUSTRE program (node) consisting of the single equation Y. Then the behaviour of the LUSTRE program is equivalent to on the level of observable signals. Now preemption may be applied since, cheerfully abusing syntax, do watching s is just: This simple trick of adding control reconciles the declarative and the imperative programming styles, at least with regard to preemption of dataflows. Another facet of the interaction between dataflow and controlflow emerges in the example discussed in Section 4.2. Recall that a LUSTRE node (iData) was placed in parallel with an ESTEREL module (iCtrl) in order to realise the initialisation phase of the production cell. We have already seen how to terminate the LUSTRE node when done is emitted, but how is the OutPress1 signal (say) to be handled by the ESTEREL control program? The point is that OutPress1 from Section 4.2.2, is a Boolean flow which is (modulo preemption) always present with a Boolean value; it is not a pure signal as required by 7- 89;:=<?>A@ B0CDBE/F0GH s in Section 4.2.1. The difficulty is that LUSTRE does not distinguish between 'clocks' and Boolean flows (although SIGNAL and ESTEREL do, for example). For strict formality we are therefore obliged to sample the Boolean flows: s s - true OutPress1 s where true is the natural projection: true z -D|V- z tt - . Ultimately, of course, the control expressions will be implemented by Boolean functions in the target (program- ming) language to which the synchronous automaton is compiled; true is thus a detail required for the consistency of the model, rather than an implementation detail. This closes our discussion of the mathematical underpinnings of the SYNCHRONIE work- bench. Of course, many technical details (of the formal translations) have been omitted in this semantics in a nutshell, but we hope that the reader will have grasped the basic, rather simple ideas. However, one important issue in synchronous programming languages that we have not dwelt upon, but which deserves a mention at least, is that of guaranteeing determinism and reactivity. The formal translations may well generate clauses such as u- @ which are inconsistent, or allow non-deterministic behaviour. Compilers can report such mishaps and indicate what went wrong. If a synchronous automaton is free of such problems (this comes under the name of causality analysis, or clock calculus), it is guaranteed to specify a deterministic Mealy machine which may be implemented either in hardware (for Boolean automata), or software. 6. The Architecture of SYNCHRONIE The SYNCHRONIE project at GMD [33] is organised around the construction a workbench for mixed synchronous language programming which combines editors, compilers, code generators and tools for simulation as well as tools for testing, verification and animation. The integration of languages is based on the synchronous automata sketched above; these play a pivotal role in the system architecture of SYNCHRONIE (see Figure 7 on page 21) since all back-end tools are based on this common representation of synchronous programs. An essential tool that is intrinsic to the workbench is the linker which merges synchronous automata from various sources; this approach guarantees a maximum of independence with regard to specific language constructs, and improves modularity in translating programs to synchronous automata. The other kernel function is a causality analyser which implements sound heuristics that guarantee reactivity and determinism-the algorithm implemented is similar to that of Shiple et al. [35]. Editor Lustre Translator Lustre Editor Translator Esterel Argos Editor Translator Argos Esterel Embedded Eiffel Embedded Eiffel Synchronie Kernel Causality Analyser Linker Synchronous Automata Validation Testing VIS Verifiers Timing Analysis Netlist Optimisers Code Generators Compilation Simulation Esterel Argos Lustre Viewers Simulator Animator Project Management Figure 7. The SYNCHRONIE workbench Project Management The increasing trend in electro-technical systems of hardware components being replaced by software provides the advantage that a device can be easily adjusted to an individual customer's needs. In this enterprise, synchronous languages themselves address the more delicate parts of the design problem (managing the often convoluted control flow, and synchronisations between distributed elements or control inputs), but this has to be recognised as only a part of the software problem. Overall, the appropriate flexibility in software design is best achieved using object oriented techniques. Although there is no space to elaborate here, this is an important part of the SYNCHRONIE toolset and we are developing an object oriented design environment for real-time applications based on a fusion of ARGOS and ESTEREL (for programming the control flow), and the object oriented language Eiffel (for the arithmetic, and other data manipulations). We call the hybrid language Embedded Eiffel. Thus, a project may have multiple components written in a variety of synchronous languages (but 'hosted' only in C or Embedded Eiffel at present). This most visible layer of the SYNCHRONIE workbench provides programmers with editing, browsing, and project management facilities amongst components written in different languages, and a control front-end to other functions of SYNCHRONIE. Compilation The compilation functions provide the link towards executing synchronous automata on various platforms. In particular, this function produces code for simulation and analysis tools for rapid prototyping and design validation. This encompasses: Optimisers may reorganise generated code variously, but the optimisations performed depend on the ultimate destination of the code. For instance, optimisations that eliminate internal signals should be inhibited when code for simulators and symbolic debuggers is required. However, experience indicates that optimisation to minimise the number of registers is desirable for formal verification tools. Code Generators exploit the close relationship between Boolean automata and sequential circuits which makes it straightforward to generate netlist formats to be exported to hardware synthesis and analysis tools such as SIS [5]. Generating portable executable code like C exploits a hierarchical representation of synchronous automata for runtime efficiency of the resulting code. Ultimately we aim to produce code for particular target architectures, like the PIC processors for micro-controller applications, but optimisation has to be applied with care, particularly in safety critical applications, as there is often a requirement for readable code (meaning that requirements can be traced to the executable code). Validation Analysis functions are intended to improve confidence throughout the steps of the development chain, helping designers and programmers to reduce the cost of errors by finding them as early as possible. Several types of analysis function are identified with Verifiers support design validation through formal proof of logical specifications and re- quirements. For the most part we provide links and interfaces to third party tools. For instance, Boolean automata equate with sequential circuits so we can use (among many adequate alternatives) the SMV model checker, or the NP-TOOLS propositional logic verifier, on aspects of the design [28, 36]. Testing improves product confidence through simulation. The testing tool allows selection between several criteria like path or boundary values testing. A prototype test specification environment to support systematic testing in the sense of [30] has been designed around the workbench's graphical ARGOS editor. This second type of verification tool complements the formal verification tools in that it focuses more on behavioural specification [31]. Synchronous automata, and particularly Boolean automata, have a simple execution model that supports fine grained timing analysis. This is important for verifying that a program can meet strong timing constraints (i.e., to satisfy the synchrony hypothesis). Simulation The simulation functions provide stepwise interpretation of synchronous au- tomata, and a means to effect rapid prototyping. Two simulation tools are distinguished: Simulators illustrate execution of synchronous programs at the source level by highlighting the syntactical entities of programs which correspond to active signals and registers. Source viewers, like editors, provide browsing facilities among the different components of a program. Animators support rapid prototyping of simple environments for executing programs. Animators are based on a toolbox of basic components for generating inputs and displaying outputs (via waveform diagrams, etc. The production cell control program described in Section 4 has been developed entirely using prototype tools from the SYNCHRONIE toolset (the figures were created using the ARGOS editor). The program itself is moderately complex, having several hundred signals and 140 control registers, but is easily handled by the VIS (a CTL [9] model checker and Verilog simulation and hardware synthesis environment [6]) for which we have a simple back-end code generator. 7. Conclusions and Further Work Embedded Eiffel has already been used to develop two small-scale industrial products, both now marketed in Germany: the first is a Massflow Meter (a sensitive coriolis device for measuring fluid flow [8]); the second is a small electronic lock system based on Radio Frequency technology [7]. Both were developed using SYNCHRONIE tools, but the high-level (ESTEREL) specification had to be hand translated into assembler in the latter case since the target technology (a PIC16C86 micro-controller with a meager 2k ROM)was not accessible through automatically generated (C) code. Automatically producing code for such limited hardware from high-level specifications is a formidable challenge that represents a next design goal of the SYNCHRONIE project. The SYNCHRONIE project is a member of the (European) Eureka project SYNCHRON: Bringing the Synchronous Technology for Real-Time Design [1]. This is dedicated to promoting synchronous programming languages in industry. The SYNCHRONIE workbench is not therefore being developed in isolation from the current generation of synchronous programming environments (which promote the respective languages independently). Links to export synchronous automata to existing tools, such as model checkers like LESAR [14], and industrial strength verification tools like NP-TOOLS [36], will emerge in due course since we plan to be fully compatible with the DC (for 'declarative code') format [10], principal deliverable of the SYNCHRON project. The idea of integrating synchronous programming languages is not entirely without precedent therefore, although the efforts to date have mainly focused on providing a common exchange format between downstream analysis tools (Halbwachs gives a good summary of the 'common formats' predating DC [12]). Jourdan and others proposed a semantical integration of ARGOS and LUSTRE, based on a translation of ARGOS into LUSTRE [20], but that first attempt was flawed in that some causally correct ARGOS programs would produce causally incorrect LUSTRE. More recently Maraninchi and Halbwachs have shown how to encode ARGOS in DC [27], and this offers a robust method of merging these two languages. The 'declarative code', as its name suggests, has been largely influenced by the dataflow synchronous programming community, and there is as yet some doubt as to its suitability for (full) ESTEREL. Synchronous automata, in contrast, supply a uniform mathematical framework through which we are able to freely intermix declarative and imperative, textual and graphical programming styles. Acknowledgments The work reported in this paper is not that of the named authors only. We should like to take this opportunity to thank all our colleagues in the Embedded Software Design group at GMD, and give them credit for their efforts in guaranteeing the success of SYNCHRONIE. The group has been led throughout by Axel Poign-e, with significant input from Reinhard Budde and Leszek Holenderski. Karl-Heinz Sylla (jointly with Budde) designed Embedded Eiffel as a paradigmatic example for the combination of object-oriented design techniques with synchronous programming; their several industrial case studies in the design of real-time systems provided insights from which the whole group has profited. Maciej Kubiczek, working closely with Leszek Holenderski, very professionally wrote compilers for ARGOS and LUSTRE, and all the supporting software in the present version of the work- bench. Agathe Merceron did many case studies in verification using the prototype tools being developed for the workbench. Monika M-ullerburg takes responsibility for all activities related to testing. Hans Nieters presented the first full version of the presentation layer including an ARGOS editor in Spring 1996 which put all our ideas to work. --R SYNCHRON: a project proposal. Synchronous languages provide safety in reactive system design. Preemption in concurrent systems. The ESTEREL synchronous programming language: Design VIS: A system for verification and synthesis. Eingebettete Echtzeitsysteme. Design and verification of synchronizing skeletons using branching time temporal logic. The common formats of synchronous languages: The declarative code DC. Specification and Design of Embedded Systems. Synchronous Programming of Reactive Systems. The synchronous data flow programming language LUSTRE. Programming and verifying real-time systems by means of the synchronous data-flow language LUSTRE visual formalism for complex systems. Production cell in LUSTRE. Software for computers in the application of industrial safety-related systems A formal approach to reactive systems software: A telecommunications application in ESTEREL. A multiparadigm language for reactive systems. Architectural synthesis for medium and high throughput processing with the New CATHEDRAL environment. Programming real-time applications with SIG- NAL Formal Development of Reactive Systems Synchronous automata for reactive Analysis of cyclic combinational circuits. Operational and compositional semantics of synchronous automaton compositions. Compiling argos into boolean equations. Symbolic Model Checking. Interpreting one concurrent calculus in another. Systematic testing: a means for validating reactive systems. Systematic testing and formal verification to validate reactive systems. Specification of complex systems. Boolean automata for implementing ESTEREL. Constructive analysis of cyclic circuits. Modelling and verifying systems and software in propositional logic. --TR --CTR Per Bjurus , Axel Jantsch, Modeling of mixed control and dataflow system in MASCOT, IEEE Transactions on Very Large Scale Integration (VLSI) Systems, v.9 n.5, p.690-703, October 2001 Axel Jantsch , Per Bjurus, Composite signal flow: a computational model combining events, sampled streams, and vectors, Proceedings of the conference on Design, automation and test in Europe, p.154-160, March 27-30, 2000, Paris, France Klaus Winkelmann, Formal Methods in Designing Embedded Systemsthe SACRES Experience, Formal Methods in System Design, v.19 n.1, p.81-110, July 2001

embedded software;synchronous automata;reactive systems;synchronous programming

290092

Filters, Random Fields and Maximum Entropy (FRAME).

This article presents a statistical theory for texture modeling. This theory combines filtering theory and Markov random field modeling through the maximum entropy principle, and interprets and clarifies many previous concepts and methods for texture analysis and synthesis from a unified point of view. Our theory characterizes the ensemble of images I with the same texture appearance by a probability distribution f(I) on a random field, and the objective of texture modeling is to make inference about f(I), given a set of observed texture examples.In our theory, texture modeling consists of two steps. (1) A set of filters is selected from a general filter bank to capture features of the texture, these filters are applied to observed texture images, and the histograms of the filtered images are extracted. These histograms are estimates of the marginal distributions of f( I). This step is called feature extraction. (2) The maximum entropy principle is employed to derive a distribution p(I), which is restricted to have the same marginal distributions as those in (1). This p(I) is considered as an estimate of f( I). This step is called feature fusion. A stepwise algorithm is proposed to choose filters from a general filter bank. The resulting model, called fields And Maximum Entropy), is a Markov random field (MRF) model, but with a much enriched vocabulary and hence much stronger descriptive ability than the previous MRF models used for texture modeling. Gibbs sampler is adopted to synthesize texture images by drawing typical samples from p(I), thus the model is verified by seeing whether the synthesized texture images have similar visual appearances to the texture images being modeled. Experiments on a variety of 1D and 2D textures are described to illustrate our theory and to show the performance of our algorithms. These experiments demonstrate that many textures which are previously considered as from different categories can be modeled and synthesized in a common framework.

Introduction Texture is an important characteristic of the appearance of objects in natural scenes, and is a powerful cue in visual perception. It plays an important role in computer vision, graphics, and image encoding. Understanding texture is an essential part of understanding human vision. Texture analysis and synthesis has been an active research area during the past three decades, and a large number of methods have been proposed, with different objectives or assumptions about the underlying texture formation processes. For example, in computer graphics, reaction-diffusion equations (Witkin and Kass 1991) have been adopted to simulate some chemical processes that may generate textures on skin of animals. In computer vision and psychology, however, instead of modeling specific texture formation process, the goal is to search for a general model which should be able to describe a wide variety of textures in a common framework, and which should also be consistent with the psychophysical and physiological understanding of human texture perception. The first general texture model was proposed by Julesz in the 1960's. Julesz suggested that texture perception might be explained by extracting the so-called 'k-th order' statis- tics, i.e., the co-occurrence statistics for intensities at k-tuples of pixels (Julesz 1962). Indeed, early works on texture modeling were mainly driven by this conjecture (Haralick 1979). A key drawback for this model is that the amount of data contained in the k-th order statistics is gigantic and thus very hard to handle when k ? 2. On the other hand, psychophysical experiments show that the human visual system does extract at least some statistics of order higher than two (Diaconis and Freeman 1981). More recent work on texture mainly focus on the following two well-established areas. One is filtering theory, which was inspired by the multi-channel filtering mechanism discovered and generally accepted in neurophysiology (Silverman et al. 1989). This mechanism suggests that visual system decomposes the retinal image into a set of sub-bands, which are computed by convolving the image with a bank of linear filters followed by some nonlinear procedures. The filtering theory developed along this direction includes the Gabor filters (Gabor 1946, Daugman 1985) and wavelet pyramids (Mallat 1989, Simoncelli etc. 1992, Coifman and Wickerhauser 1992, Donoho and Johnstone 1994). The filtering methods show excellent performance in classification and segmentation (Jain and Farrokhsia 1991). The second area is statistical modeling, which characterizes texture images as arising from probability distributions on random fields. These include time series models (Mc- Cormick and Jayaramamurthy 1974), Markov chain models (Qian and Terrington 1991), and Markov random field (MRF) models (Cross and Jain 1983, Mao and Jain 1992, Yuan and Rao 1993). These modeling approaches involve only a small number of parameters, thus provide concise representation for textures. More importantly, they pose texture analysis as a well-defined statistical inference problem. The statistical theories enable us not only to make inference about the parameters of the underlying probability models based on observed texture images, but also to synthesize texture images by sampling from these probability models. Therefore, it provides a rigorous way to test the model by checking whether the synthesized images have similar visual appearances to the textures being modeled (Cross and Jain 1983). But usually these models are of very limited forms, hence suffer from the lack of expressive power. This paper proposes a modeling methodology which is built on and directly combines the above two important themes for texture modeling. Our theory characterizes the ensemble of images I with the same texture appearances by a probability distribution f(I) on a random field. Then given a set of observed texture examples, our goal is to infer f(I). The derivation of our model consists of two steps. (I) A set of filters is selected from a general filter bank to capture features of the texture. The filters are designed to capture whatever features might be thought to be characteristic of the given texture. They can be of any size, linear or nonlinear. These filters are applied to the observed texture images, and histograms of the filtered images are extracted. These histograms estimate the marginal distributions of f(I). This step is called feature extraction. (II) Then a maximum entropy distribution p(I) is constructed, which is restricted to match the marginal distributions of f(I) estimated in step (I). This step is called feature fusion. A stepwise algorithm is proposed to select filters from a general filter bank, and at each step k it chooses a filter F (k) so that the marginal distributions of f(I) and p(I) with respect to F (k) have the biggest distance in terms of L 1 norm. The resulting model, called fields And Maximum Entropy), is a Markov random field (MRF) model 1 , but with a much more enriched vocabulary and hence much stronger descriptive power compared with previous MRF models. The Gibbs sampler is adopted to synthesize texture images by drawing samples from p(I), thus the model is tested by synthesizing textures in both 1D and 2D experiments. Our theory is motivated by two aspects. Firstly, a theorem proven in section (3.2) shows that a distribution f(I) is uniquely determined by its marginals. Therefore if a model p(I) matches all the marginals of f(I), then recent psychophysical research on human texture perception suggests that two 'homogeneous' textures are often difficult to discriminate when they have similar marginal distributions from a bank of filters (Bergen and Adelson 1991, Chubb and Landy 1991). Our method is inspired by and bears some similarities to Heeger and Bergen's (1995) recent work on texture synthesis, where many natural looking texture images were synthesized by matching the histograms of filter responses organized in the form of a pyramid. This paper is arranged as follows. First we set the scene by discussing filtering methods and Markov random field models in section (2), where both the advantages and disadvantages of these approaches are addressed. Then in section (3), we derive our FRAME model and propose a feature matching algorithm for probability inference and stochastic simulation. Section (4) is dedicated to the design and selection of filters. The texture modeling experiments are divided into three parts. Firstly section (5) illustrates important concepts of the FRAME model by designing three experiments for one dimensional texture synthesis. Secondly a variety of 2D textures are studied in section (6). Then section (7) discusses a special diffusion strategy for modeling some typical texton images. Finally section (8) concludes with a discussion and the future work. Among statisticians, MRF usually refers to those models where the Markov neighborhood is very small, e.g. 2 or 3 pixels away. Here we use it for any size of neighborhood. Filtering and MRF Modeling 2.1 Filtering theory In the various stages along the visual pathway, from retina, to V1, to extra-striate cortex, cells with increasing sophistication and abstraction have been discovered: center-surround isotropic retinal ganglion cells, frequency and orientation selective simple cells, and complex cells that perform nonlinear operations. Motivated by such physiological discoveries, the filtering theory proposes that the visual system decomposes a retinal image into a set of sub-band images by convolving it with a bank of frequency and orientation selective linear filters. This linear filtering process is then followed by some nonlinear operations. In the design of various filters, Gaussian function plays an important role due to its nice low-pass frequency property. To fix notation, we define an elongated two-dimensional Gaussian function as: with location parameters parameters (oe x ; oe y ). A simple model for the radially symmetric center-surround ganglion cells is the Laplacian of Gaussian with oe Similarly, a model for the simple cells is the Gabor filter (Daugman 1985), which is a pair of cosine and sine waves with frequency ! and amplitude modulated by the Gaussian By carefully choosing the frequency ! and rotating the filter in the x-y coordinate system, we obtain a bank of filters which cover the entire frequency domain. Such filters are used for image analysis and synthesis successfully by (Jain and Farrokhsia 199, Lee 1992 ). Other filter banks have also been designed for image processing by (Simoncelli etc. 1992). The filters mentioned above are linear. Some functions are further applied to these linear filters to model the nonlinear functions of the complex cell. One way to model the complex cell is to use the power of each pair of Gabor filter j In fact, is the local spectrum S(!) of I at (x; y) smoothed by a Gaussian function. Thus it serves as a spectrum analyzer. Although these filters are very efficient in capturing local spatial features, some problems are not well understood. For example, i) given a bank of filters, how to choose the best set of filters? especially when some of the filters are linear while others are nonlinear, or the filters are highly correlated to each other, ii) after selecting the filters, how to fuse the features captured by them into a single texture model? These questions will be answered in the rest of the paper. 2.2 MRF modeling MRF models were popularized by Besag (Besag 1973) for modeling spatial interactions on lattice systems and were used by (Cross and Jain 1983) for texture modeling. An important characteristic of MRF modeling is that the global patterns are formed via stochastic propagation of local interactions, which is particularly appropriate for modeling textures since they are characterized by global but not predictable repetitions of similar local structures. In MRF models, a texture is considered as a realization from a random field I defined over a spatial configuration D, for example, D can be an array or a lattice. We denote I(~v) as the random variable at a location ~v 2 D, and let Dg be a neighborhood system of D, which is a collection of subsets of D satisfying 1) . The pixels in N ~v are called neighbors of ~v. A subset C of D is a clique if every pair of distinct pixels in C are neighbors of each other; C denotes the set of all cliques. Definition. p(I) is an MRF distribution with respect to N if p(I(~v) j denotes the values of all pixels other than ~v, and for A ae D, denotes the values of all pixels in A. Definition. p(I) is a Gibbs distribution with respect to N if Z where Z is the normalizing constant (or partition function), and -C () is a function of intensities of pixels in clique C (called potential of C). Some constraints can be imposed on -C for them to be uniquely determined. The Hammersley-Clifford theorem establishes the equivalence between MRF and the Gibbs distribution (Besag Theorem 1 For a given N , p(I) is an MRF distribution () p(I) is a Gibbs distribution. This equivalence provides a general method for specifying an MRF on D, i.e., first choose an N , and then specify -C . The MRF is stationary if for every C 2 C, -C depends only on the relative positions of its pixels. This is often assumed in texture modeling. Existing MRF models for texture modeling are mostly auto-models (Besag 1973) with potentials, i.e. -C j 0 if jCj ? 2, and p(I) has the following form Z expf where and ~v are neighbors. The above MRF model is usually specified through conditional distributions, where the neighborhood is usually of order less than or equal to three pixels, and some further restrictions are usually imposed on g for p(I(~v) j I(\Gamma~v)) to be a linear regression or the generalized linear model. Two commonly used auto-models are the auto-binomial model and the auto-normal model. The auto-binomial model is used for images with finite grey levels I(~v) 2 f0; 1g (Cross and Jain 1983), the conditional distribution is a logistic regression, where log p ~v It can be shown that the joint distribution is of the form Z expf G When 2, the auto-binomial model reduces to the auto-logistic model (i.e., the Ising model), which is used to model binary images. The auto-normal model is used for images with continuous grey levels (Yuan and Rao 1993). It is evident that if an MRF p(I) is a multivariate normal distribution, then p(I) must be auto-normal, so the auto-normal model is also called a Gaussian MRF or GMRF. The conditional distribution is a normal regression, and p(I) is of the form i.e., the multivariate normal distribution N(-; oe the diagonal elements of B are unity and the off-diagonal (~u; ~v) element of it is \Gammafi ~u\Gamma~v . Another MRF model for texture is the OE-model (Geman and Graffigne 1986): Z expf\Gamma where OE is a known even symmetric function, and the OE-model can be viewed as extended from the Potts model (Winkler 1995). The advantage of the auto-models is that the parameters in the models can be easily inferred by auto-regression, but they are severely limited in the following two aspects: i) the cliques are too small to capture features of texture, ii) the statistics on the cliques specifies only the first-order and second order moments (e.g. means and covariances for GMRF). However, many textures has local structures much larger than three or four pixels, and the covariance information or equivalently spectrum can not adequately characterize textures, as suggested the existence of distinguishable texture pairs with identical second-order or even third-order moments, as well as indistinguishable texture pairs with different second-order moments (Diaconis and Freeman 1981). Moreover many textures are strongly non-Gaussian, regardless of neighborhood size. The underlying cause of these limitations is that equation 3 involves too many parameters if we increase the neighborhood size or the order of the statistics, even for the simplest auto-models. This suggests that we need carefully designed functional forms for -C () to efficiently characterize local interactions as well as the statistics on the local interactions. 3 From maximum entropy to FRAME model 3.1 The basics of maximum entropy Maximum entropy (ME) is an important principle in statistics for constructing a probability distributions p on a set of random variables X (Jaynes 1957). Suppose the available information is the expectations of some known functions OE R Let\Omega be the set of all probability distribution p(x) which satisfy the constraints, i.e., The ME principle suggests that a good choice of the probability distribution is the one that has the maximum entropy, i.e., Z log p(x)dxg; (11) subject to E p [OE n R and R By Lagrange multipliers, the solution for p(x) is: where is the Lagrange parameter, and Z ( R expf\Gamma is the partition function that depends on and it has the following properties: @ log Z Z In equation (12), (- 1 ; :::; -N ) is determined by the constraints in equation (11). But a closed form solution for not available in general, especially when OE n (\Delta) gets complicated, so instead we seek numerical solutions by solving the following equations iteratively. dt The second property of the partition function Z ( ) tells us that the Hessian matrix of log Z ( ) is the covariance matrix of vector (OE 1 (x); OE 2 (x); :::; OE N (x)) which is definitely positive 2 , and log Z ( ) is strictly concave with respect to (- 1 ; :::; -N ), so is log p(x; ). Therefore given a set of consistent constraints, there is a unique solution for in equation (13). 3.2 Deriving the FRAME model Let image I be defined on a discrete domain D, D can be a N \Theta N lattice. For each pixel and L is an interval of R or L ae Z. For each texture, we assume that there exists a "true" joint probability density f(I) over the image space L jDj , and f(I) should concentrate on a subspace of L jDj which corresponds to texture images that have perceptually similar texture appearances. Before we derive the FRAME model, we first fix the notation as below. Given an image I and a filter F (ff) with being an index of filter, we let I (ff) I(~v) be the filter response at location ~v, and I (ff) the filtered image. The marginal empirical distribution (histogram) of I (ff) is where ffi() is the Dirac delta function. The marginal distribution of f(I) with respect to F (ff) at location ~v is denoted by f (ff) Z Z I (ff) (~v)=z At first thought, it seems an intractable problem to estimate f(I) due to the overwhelming dimensionality of image I. To reduce dimensions, we first introduce the following theorem. Here, it is reasonable to assume that OE n (x) is independent of OE j (x) if i Theorem 2 Let f(I) be the j D j-dimensional continuous probability distribution of a texture, then f(I) is a linear combination of f (-) , the latter are the marginal distributions on the linear filter response F (-) I. [Proof ]. By inverse Fourier transform, we have Z Z f (-) is the characteristic function of f(I), and Z Z e \Gamma2- i!-; I? f(I)dI Z e \Gamma2- iz dz Z Z f(I)dI Z e \Gamma2- iz dz Z Z Z e \Gamma2- iz f (-) (z)dz is the inner product, and by definition f (-) R R is the marginal distribution of F (-) I, and we define F (-) as a linear filter. 2 Theorem 2 transforms f(I) into a linear combination of its one dimensional marginal distributions. 3 Thus it motivates a new method for inferring f(I): construct a distribution p(I) so that p(I) has the same marginal distributions f (-) . If p(I) matches all marginal distributions of f(I), then f(I). But this method will involve uncountable number of filters and each filter F (-) is as big as image I. Our second motivation comes from recent psychophysical research on human texture perception, and the latter suggests that two homogeneous textures are often difficult to discriminate when they produce similar marginal distributions for responses from a bank of filters (Bergen and Adelson 1991, Chubb and Landy 1991). This means that it is plausible to ignore some statistical properties of f(I) which are not important for human texture discrimination. To make texture modeling a tractable problem, in the rest of this paper we make the following assumptions to limit the number of filters and the window size of each 3 It may help understand the spirit of this theorem by comparing it to the slice-reconstruction of 3D volume in tomography. filter for computational reason, though these assumptions are not necessary conditions for our theory to hold true. 1). We limit our model to homogeneous textures, thus f(I) is stationary with respect to location ~v. 4 2). For a given texture, all features which concern texture perception can be captured by "locally" supported filters. In other words, the sizes of filters should be smaller than the size of the image. For example, the size of image is 256 \Theta 256 pixels, and the sizes of filters we used are limited to be less than 33 \Theta 33 pixels. These filters can be linear or non-linear as we discussed in section (2.1). 3). Only a finite set of filters are used to estimate f(I) Assumptions 1) and 2) are made because we often have access to only one observed (training) texture image. For a given observed image I obs and a filter F (ff) , we let I obs(ff) denote the filtered image, and H obs(ff) (z) the histogram of I obs(ff) . According to assumption 1), f (ff) is independent of ~v. By ergodicity, H obs(ff) (z) makes a consistent estimator to f (ff) (z). Assumption 2) ensures that the image size is lager relative to the support of filters, so that ergodicity takes effect for H obs(ff) (z) to be an accurate estimate of f (ff) (z). Now for a specific texture, let Kg be a finite set of well selected filters, and f (ff) (z); are the corresponding marginal distributions of f(I). We denote the probability distribution p(I) which matches these marginal distributions as a set, is the marginal distribution of p(I) with respect to filter F (ff) at location ~v. Thus according to assumption 3), any p(I) 2\Omega is perceptually a good enough model for the texture, provided that we have enough well selected filters. Then we choose from\Omega a distribution p(I) which has the maximum entropy, Z p(I) log p(I)dIg; (15) subject to E and R 4 Throughout this paper, we use circulant boundary conditions. The reason for us to choose the maximum entropy (ME) distribution is that while p(I) satisfies the constraints along some dimensions, it is made as random as possible in other unconstrained dimensions, since entropy is a measure of randomness. In other words, p(I) should represent information no more than that is available. Therefore an ME distribution gives the simplest explanation for the constraints and thus the purest fusion of the extracted features. The constraints on equation (15) differ from the ones given in section (3.1) in that z takes continuous real values, hence there are uncountable number of constraints, therefore, the Lagrange parameter - takes the form as a function of z. Also since the constraints are the same for all locations ~v 2 D, - should be independent of ~v. Solving this maximization problem gives the ME distribution: Z is a set of selected filters, and is the Lagrange parameter. Note that in equation (17), for each filter F (ff) , - (ff) () takes the form as a continuous function of the filter response I (ff) (~v). To proceed further, let's derive a discrete form of equation (17). Assume that the filter response I (ff) (~v) is quantitized into L discrete grey levels, therefore z takes values from set fz (ff) L g. In general, the width of these bins do not have to be equal, and the number of grey levels L for each filter response may vary. As a result, the marginal distributions and histograms are approximated by piecewisely constant functions of L bins, and we denote these piecewise functions as vectors. H L ) is the histogram of I (ff) , H obs(ff) denotes the histogram of I obs(ff) , and the potential function - (ff) () is approximated by vector - equation (16) is rewritten as: by changing the order of summations: The virtue of equation (18) is that it provides us with a simple parametric model for the probability distribution on I, and this model has the following properties, specified by ffl Given an image I, its histograms H (1) are sufficient statistics, i.e. p(I; K ; SK ) is a function of (H We plug equation (18) into the constraints of the ME distribution, and solve for iteratively by the following equations, d- (ff) dt In equation (19), we have substituted H obs(ff) for f (ff) , and E p(I; K ;S K ) (H (ff) ) is the expected histogram of the filtered image I (ff) where I follows p(I; K ; SK ) with the current K . Equation (19) converges to the unique solution at K as we discussed in section (3.1), and - K is called the ME-estimator. It is worth mentioning that this ME-estimator is equivalent to the maximum likelihood estimator (MLE), log p(I obs ; K log Z ( K By gradient accent, maximizing the log-likelihood gives equation (19), following property i) of the partition function Z ( K ). In Equation (19), at each step, given K and hence p(I; K ; SK ), the analytic form of available, instead we propose the following method to estimate it as we did for f (ff) before. We draw a typical sample from p(I; K ; SK ), and thus synthesize a texture image I syn . Then we use the histogram H syn(ff) of I syn(ff) to approximate This requires that the size of I syn that we are synthesizing should be large enough. 5 To draw a typical sample image from p(I; K ; SK ), we use the Gibbs sampler (Geman and Geman 1984) which simulates a Markov chain in the image space L jDj . The Markov chain starts from any random image, for example, a white noise image, and it converges to a stationary process with distribution p(I; K ; SK ). Thus when the Gibbs sampler converges, the images synthesized follow distribution p(I; K ; SK ). In summary, we propose the following algorithm for inferring the underlying probability model p(I; K ; SK ) and for synthesizing the texture according to p(I; K ; SK ). The algorithm stops when the subband histograms of the synthesized texture closely match the corresponding histograms of the observed images. 6 5 Empirically, 128 \Theta 128 or 256 \Theta 256 seems to give a good estimation. 6 We assume the histogram of each subband I (ff) is normalized such that all the computed in this algorithm have one extra degree of freedom for each ff, i.e., we can increase f- (ff) by a constant without changing p(I; K ; SK ). This constant will be absorbed by the partition function Z ( K ). Algorithm 1. The FRAME algorithm Input a texture image I obs . Select a group of K filters g. Compute Kg. Initialize - (ff) Initialize I syn as a uniform white noise texture. 7 Repeat Calculate H syn(ff) :::; K from I syn , use it for E p(I; K ;SK Update - (ff) Apply Gibbs sampler to flip I syn for w sweeps under p(I; K Until 1P L Algorithm 2. The Gibbs Sampler for w sweeps Given image I(~v), flip counter/ 0 Repeat Randomly pick a location ~v under the uniform distribution. For being the number of grey levels of I Calculate Randomly flip I(~v) / val under p(val j I(\Gamma~v)). flip counter / flip counter Until flip counter=w \Theta M \Theta N . In algorithm 2, to compute I(\Gamma~v)), we set I(~v) to val, due to Markov property, we only need to compute the changes of I (ff) at the neighborhood of ~v. The size of the neighborhood is determined by the size of filter F (ff) . With the updated I (ff) , we calculate H (ff) , and the probability is normalized such that 1. In the Gibbs sampler, flipping a pixel is a step of the Markov chain, and we define flipping jDj pixels as a sweep, where jDj is the size of the synthesized image. Then the overall iterative process becomes an inhomogeneous Markov chain. At the beginning of the process, p(I; K ; SK ) is a "hot" uniform distribution. By updating the parameters, the process get closer and closer to the target distribution, which is much colder. So the algorithm is very much like a simulated annealing algorithm (Geyer and Thompson 1995), 7 Note that the white noise image with uniform distribution are the samples from p(I; K ; SK ) with which is helpful for getting around local modes of the target distribution. We refer to (Winkler 1995) for discussion of alternative sampling methods. The computational complexity of the above algorithm is notoriously large: O(U \Theta w \Theta jDj \Theta G \Theta K \Theta FH \Theta FW ) with U the number of updating steps for K , w the number of sweeps each time we update K , D the size of image I syn , G the number of grey levels of I, K the number of filters, and FH;FW are the average window sizes of the filters. To synthesize a 128 \Theta 128 texture, the typical complexity is about 50 \Theta 4 \Theta 128 \Theta 128 \Theta 8 \Theta 4 \Theta takes about one day to run on a Sun-20. Therefore it is very important to choose a small set of filter which can best capture the features of the texture. We discuss how to choose filters in section (4). 3.3 A general framework The probability distribution we derived in the last section is of the form This model is significant in the following aspects. First, the model is directly built on the features I (ff) (~v) extracted by a set of filters F (ff) . By choosing the filters, it can easily capture the properties of the texture at multiple scales and orientations, either linear or nonlinear. Hence it is much more expressive than the cliques used in the traditional MRF models. Second, instead of characterizing only the first and second order moments of the marginal distributions as the auto-regression MRF models did, the FRAME model includes the whole marginal distribution. Indeed, in a simplified case, if the constraints on the probability distribution are given in the form of kth-order moments instead of marginal distributions, then the functions - (ff) (\Delta) in equation (21) become polynomials of order m. In such case, the complexity of the FRAME model is measured by the following two aspects: 1) the number of filters and the size of the filter, 2) the order of the moments-m. As we will see in later experiments, equation (21) enable us to model strongly non-Gaussian textures. It is also clear to us that all existing MRF texture models can be shown as special cases of FRAME models. Finally, if we relax the homogeneous assumption, i.e., let the marginal distribution of I (ff) (~v) depend on ~v, then by specifying these marginal distributions, denoted by f (ff) p(I) will have the form Z expf\Gamma This distribution is relevant in texture segmentation where - (ff) ~v are assumed piecewise consistent with respect to ~v, and in shape inference when - (ff) ~v changes systematically with respect to ~v, resulting in a slowly varying texture. We shall not study non-stationary textures in this paper. In summary, the FRAME model incorporates and generalizes the attractive properties of the filtering theory and the random fields models, and it interprets many previous methods for texture modeling in a unified view of point. 4 Filter selection In the last section, we build a probability model for a given texture based on a set of filters SK . For computational reasons SK should be chosen as small as possible, and a key factor for successful texture modeling is to choose the right set of filters that best characterizes the features of the texture being modeled. In this section, we propose a novel method for filter selection. 4.1 Design of the filter bank To describe a wide variety of textures, we first need to design a filter bank B. B can include all previously designed multi-scale filters (Daugman 1985, Simoncelli etc 1992) or wavelets (Mallat 1989, Donoho and Johnstone 1994, Coifman and Wickerhauser 1992). In this paper, we should not discuss the optimal criterion for constructing a filter bank. Throughout the experiments in this paper, we use five kinds of filters. 1. The intensity filter ffi(), and it captures the DC component. 2. The isotropic center-surround filters, i.e., the Laplacian of Gaussian filters. Here we rewrite equation (1) as: const 2oe stands for the scale of the filter. We choose eight scales with these filters by LG(T ). 3. The Gabor filters with both sine and cosine components. We choose a simple case from equation (2): const We choose 6 scales We can see that these filters are not even approximately orthogonal to each other. We denote by Gcos(T ; ') and Gsin(T ; ') the cosine and sine components of the Gabor filters. 4. The spectrum analyzers denoted by SP (T ; '), whose responses are the power of the Gabor pairs: j (Gabor I)(x; y) j 2 . 5. Some specially designed filters for one dimensional textures and the textons, see section (5) and (7). 4.2 A stepwise algorithm for filter selection For a fixed model complexity K, one way to choose SK from B is to search for all possible combinations of K filters in B and compute the corresponding p(I; K ; SK ). Then by comparing the synthesized texture images following each p(I; K ; SK ), we can see which set of filters is the best. Such a brute force search is computationally infeasible, and for a specific texture we often do not know what K is. Instead, we propose a stepwise greedy strategy. We start from S an uniform distribution, and then sequentially introduce one filter at a time. Suppose that at the k-th step we have chosen S obtained a maximum entropy distribution so that E p(I; k ;S k ) [H (ff) k. Then at the (k 1)-th step, for each filter F (fi) 2 B=S k , we denote by the distance between are respectively the marginal distributions of p(I; k and f(I) with respect to filter F (fi) . Intuitively, the bigger d(fi) is, the more information F (fi) carries, since it reports a big difference between p(I; k ; S k ) and f(I). Therefore we should choose the filter which has the maximal distance, i.e., Empirically we choose to measure the distance d(fi) in terms of L p -norm, i.e., In the experiments of this paper, we choose To estimate f (fi) and E p(I; k ;S k ) [H (fi) ], we applied F (fi) to the observed image I obs and the synthesized image I syn sampled from the p(I; k ; S k ), and substitute the histograms of the filtered images for f (fi) and E p(I; k ;S k ) [H (fi) ], i.e., The filter selection procedure stops when the d(fi) for all filters F (fi) in the filter bank are smaller than a constant ffl. Due to fluctuation, various patches of the same observed texture image often have a certain amount of histogram variance, and we use such a variance for ffl. In summary, we propose the following algorithm for filter selection. Algorithm 3. Filter Selection Let B be a bank of filters, S the set of selected filters, I obs the observed texture image, and I syn the synthesized texture image. dist. I syn / uniform noise. For do Compute I obs(ff) by applying F (ff) to I obs . Compute of I obs(ff) . Repeat For each F (fi) 2 B=S do Compute I syn(fi) by applying F (fi) to I syn Compute of I syn(fi) Choose F (k+1) so that d(k S Starting from p(I) and I syn , run algorithm 1 to compute new p (I) and I syn . p(I) / p (I) and I syn / I syn . Until d (fi) ! ffl Before we conclude this section, we would like to mention that the above criterion for filter selection is related to the minimax entropy principle studied in (Zhu, Wu, and Mumford 1996). The minimax entropy principle suggests that the optimal set of filters SK should be chosen to minimize the Kullback-Leibler distance between p(I; K ; SK ) and f(I), and the latter is measured by the entropy of the model p(I; K ; SK ) up to a constant. Thus at each step k filter is selected so that it minimizes the entropy of p(I; k by gradient descent, i.e., and + is the new Lagrange parameter. A brief derivation is given in the appendix. 5 Experiments in one dimension In this section we illustrate some important concepts of the FRAME model by studying a few typical examples for 1D texture modeling. In these experiments, the filters are chosen by hand. For one-dimensional texture the domain is a discrete array and a pixel is indexed by x instead of ~v. For any x 2 [0; 255], I(x) is discretized into G grey levels, with experiment I and III, and experiment II. Experiment I. This experiment is designed to show the analogy between filters in texture modeling and vocabulary in language description, and to demonstrate how a both histograms are normalized to have We note this measure is robust with respect to the choice of the bin number L (e.q. we can take as well as the normalization of the filters. a Figure 1 The observed and synthesized pulse textures. a) the observed; b) synthesized using only the intensity filter; c) intensity filter plus rectangular filter with d) Gabor filter with Gabor filter plus intensity filter. texture can be specified by the marginal distributions of a few well selected filters. The observed texture, as shown in figure (1.a), is a periodic pulse signal with period once every 8 pixels and all the other pixels. First we choose an intensity filter, and the filter response is the signal itself. The synthesized texture by FRAME is displayed in figure (1.b). Obviously it has almost the same number of pulses as the observed one, and so has approximately the same marginal distribution for intensity. Unlike the observed texture, however, these pulses are not arranged periodically. To capture the period of the signal, we add one more special filter, an 8 \Theta 1 rectangular 1], and the synthesized signal is shown in figure (1.c), which has almost the same appearance as in Figure (1.a). We can say that the probability p(I) specified by these two filters models the properties of the input signal very well. Figure (1.d) is the synthesized texture using a nonlinear filter which is an 1D spectrum analyzer SP (T ) with 8. Since the original periodic signal has flat power spectrum, and the Gabor filters only extract information in one frequency band, therefore the texture synthesized under p(I) has power spectrum near frequency 2-but are totally free at other bands. Due to the maximum entropy principle, the FRAME model allows for the unconstrained frequency bands to be as noisy as possible. This explains why figure (1.d) is noise like while having roughly a period of 8. If we add the intensity filter, then probability p(I) captures the observed signal again, and a synthesized texture is shown in figure (1.e). This experiment shows that the more filters we use, the closer we can match the synthesized images to the observed. But there are some disadvantages for using too many filters. Firstly it is computationally expensive, and secondly, since we have few observed examples, it may overly constrain the probability p(I), i.e. it may make p(I) 'colder' than it should be. Experiment II. In this second experiment we compare FRAME against Gaussian MRF models by showing the inadequacy of the GMRF model to express high order statistics. To begin with, we choose a gradient filter r with impulse response [\Gamma1; 1] for com- parison, and the filtered image is denoted by rI. The GMRF models are concerned with only the mean and variance of the filter re- sponses. As an example, we put the following two constraints on distribution p(I), Since we use a circulant boundary, the first constraint always holds, and the resulting maximum entropy probability is Z expf\Gamma- The numeric solution given by the FRAME algorithm is and two synthesized texture images are shown in figure (3.b) and (3.c). Figure (3.a) is a white noise texture for comparison. x x I x x I x x . . . 94 a b Figure a) The designed marginal distribution of rI; and b) The designed marginal distribution of \DeltaI. As a comparison, we now ask rI(x) to follow the distribution shown in figure (2.a). Clearly in this case E p [rI(x)] is a non-Gaussian distribution with first and second moments as before, i.e., synthesized textures are displayed in figure (3.d and e). The textures in figure (3.d and e) possess the same first and second order moments as in figure (3.b and c), but figure (3.d and e) have specific higher order statistics and looks more specific than in figure (3.b and c). It demonstrates that the FRAME model has more expressive power than the GMRF model. Now we add a Laplacian filter \Delta with impulse response [0:5; \Gamma1:0; 0:5], and we ask \DeltaI(x) to follow the distribution shown in figure (2.b). Clearly the number of peaks and valleys in I(x) are specified by the two short peaks in figure (2.b), the synthesized texture c d Figure 3 a. The uniform white noise texture, b.c. the texture of GMRF, d,e, the texture with higher order statistics, f. the texture specified with one more filter. is displayed in figure (3.f). This experiment also shows the analogy between filters and vocabulary. Experiment III. This experiment is designed to demonstrate how a single non-linear Gabor filter is capable of forming global periodic textures. The observed texture is a perfect sine wave with period T hence it has a single Fourier component. We choose the spectrum analyzer SP (T ) with period 16. The synthesized texture is in figure (4.a). The same is done for another sine wave that has period T and correspondingly the result is shown in figure (4.b). Figure (4) show clear globally periodic signals formed by single local filters. The noise is due to the frequency resolution of the filters. Since the input textures are exactly periodic, the optimal resolution will requires the Gabor filters to be as long as the input signal, which is computationally more expensive. Figure 4 The observed textures are the pure sine waves with period T=16, and 32 respectively. Periodic texture synthesized by a pair of Gabor filters, a. T=16, b. T=32. 6 Experiments in two dimensions In this section, we discuss texture modeling experiments in two dimensions. We first take one texture as an example to show in detail the procedure of algorithm 3, then we will apply algorithm 3 to other textures. Figure (5.a) is the observed image of animal fur. We start from the uniform noise image in figure (5.b). The first filter picked by the algorithm is a Laplacian of Gaussian filter LG(1:0) and its window size is 5 \Theta 5. It has the largest error all the filters in the filters bank. Then we synthesize texture as shown in figure (5.c), which has almost the same histogram at the subband of this filter (the error d(fi) drops to 0:035). Comparing figure (5.c) with figure (5.b), we notice that this filter captures local smoothness feature of the observed texture. Then the algorithm sequentially picks 5 more filters. They are 2) Gcos(6:0; 120 each of which captures features at various scales and orientations. The sequential conditional errors for these filters are respectively 0:424; 0:207; 0:132; 0:157; 0:059 and the texture images synthesized using filters are shown in figure (5.d,e,f). Obviously, with more filters added, the synthesized texture gets closer to the observed one. To show more details, we display the subband images of the 6 filters in figure (6), the histograms of these subbands H (ff) and the corresponding estimated parameters - (ff) are plotted in figure (7) and figure (8) respectively. Figure 5 Synthesis of the fur texture. a is the observed texture, b,c,d,e,f are the synthesized textures using filters respectively. See text for interpretation. Figure 6 The subband images by applying the 6 filters to the fur image: a Laplacian of Gaussian a b c Figure 7 a,b,c,d,e,f are respectively the histograms H (ff) for lambda i lambda i a b c lambda i lambda i 3 41.53.5i lambda i Figure 8 a,b,c,d,e,f are respectively the - (ff) for In figure (7), the histograms are approximately Gaussian functions, and correspon- dently, the estimated - (ff) in figure (8) are close to quadratic functions. Hence in this example, the high order moments seemly do not play a major role, and the probability model can be made simpler. But this will not be always true for other textures. In figure (8), we also notice that the computed - (ff) becomes smaller and smaller when ff gets bigger, which suggests that the filters chosen in later steps make less and less contribution to p(I), and thus confirms our early assumption that the marginal distributions of a small number of filtered images are good enough to capture the underlying probability distribution f(I). Figure (9.a) is the scene of the mud ground with footprints of animals, these footprints are filled with water and get brighter. This is a case of sparse features. Figure (9.b) is the synthesized texture using 5 filters chosen by algorithm 3. Figure (10.a) is an image taken from the skin of cheetah. the synthesized texture using 6 filters is displayed in figure (10.b). We notice that in figure (10.a) the texture is not homogeneous, the shapes of the blobs vary with spatial locations and the left upper corner is darker than the right lower one. The synthesized texture, shown in figure (10.b), also has elongated blobs introduced by different filters, but we notice that the bright pixels Figure 9 a. the observed texture-mud, b, the synthesized one using 5 filters a b Figure a. the observed texture-cheetah blob, b, the synthesized one using 6 filters c d Figure 11 a. the input image of fabric, b. the synthesized image with two spectrum analyzers plus the Laplacian of Gaussian filter. c,d The filter response of the two spectrum analyzers for the fabric texture spread uniformly across the image. Finally we show a texture of fabric in figure (11.a), which has clear periods along both horizontal and vertical directions. We want to use this texture to test the use of non-linear filters, so we choose two spectrum analyzers to capture the first two salient periods, one in the horizontal direction, the other in the vertical direction. The filter responses I (ff) are the sum of squares of the sine and cosine component responses. The filter responses are shown in figure (11c, d), and are almost constant. We also use the intensity filter and the Laplacian of Gaussian filter LG( (with window size 3 \Theta to take care of the intensity histogram and the smoothness. The synthesized texture is displayed in figure (11.b). If we carefully look at figure (11.b), we can see that this synthesized texture has mis-arranged lines at two places, which may indicate that the sampling process was trapped in a local maximum of p(I). 7 The Sampling strategy for textons In this section, we study a special class of textures formed from identical textons, which psychophysicists studied extensively. Such texton images are considered as rising from a different mechanism from other textures in both psychology perception and previous texture modeling, and the purpose of this section is to demonstrate that they can still be modeled by the FRAME model, and to show an annealing strategy for computing Figure (12.a,b) are two binary (\Gamma1; +1 for black and white pixels) texton images with circle and cross as the primitives. These two image are simply generated by sequentially superimposing 128 15\Theta15 masks on a 256\Theta256 lattice using uniform distribution, provided that the dropping of one mask does not destroy the existing primitives. At the center of the mask is a circle (or a cross). For these textures, choosing filters seems easy: we simply select the above 15 \Theta 15 mask as the linear filter. Take the circle texton as an example. By applying the filter to the circle image and a uniform noise image, we obtain the histograms H obs (solid curve) and H(x) (dotted curve) plotted in figure (13.a). We observe that there are many isolated a b Figure Two typical texton images. a circle, b cross peaks in H obs , which set up " potential wells" so that it becomes extremely unlikely to change a filter response at a certain location from one peak to another by flipping one pixel at a time. -22Filter response lambda a b Figure a. The solid curve is the histogram of the circle image, and the dotted curve is the histogram of the noise image; b. the estimated -() function in the probability model for the image of circles To facilitate the matching process, we propose the following heuristics. We smooth H obs with a Gaussian window G oe , or equivalently run the "heat" diffusion equation on H obs (x; t) within the interval [x are respectively the minimal and maximal filter response. dH obs (x; t) dt @x @x The boundary conditions help to preserve the total "heat". Obviously, the larger t is, the smoother the H obs (x; t) will be. Therefore we start from matching H(x) to H obs (x; t) with a large t (see figure (14.a)), then gradually decrease t and match H(x) to the histograms shown in figure (14.b,c,d,e,f) sequentially. This process is similar to the simulated annealing method. The intuitive idea is to set up "bridges" between the peaks in the original histogram, which encourages the filter response change to the two ends, where the texton forms, then we gradually destruct these "bridges". At the end of the process, the estimated - function for the circle texton is shown in figure (13.b), and the synthesized images are shown in figure (15). We notice that the cross texton is more difficult to deal with because it has slightly more complex structures than the circle, and may need more carefully designed filters. Histogram Histogram Histogram a b c Histogram Histogram Histogram Figure 14 The diffused histogram H obs smaller and smaller from a to f. Figure Two synthesized texton images. Although there is a close relationship between FRAME and the previous MRF models, the underlying philosophies are quite different. Traditional MRF approaches favor the specification of conditional distributions (Besag 1973). For auto-models, p(I(~v) j I(\Gamma~v)) are linear regressions or logistic regressions, so the modeling, inference, and interpretation can be done in a traditional way. While it is computationally efficient for estimating the fi coefficients, this method actually limits our imagination for building a general model. Since the only way to generalize auto-models in the conditional distribution framework is to either increase neighborhood size, and thus introduce more explanatory variables in these auto-regressions, or introduce interaction terms (i.e., high order product terms of the explanatory variables). However, even with a modest neighborhood (e.g., 13 \Theta 13), the parameter size will be too large for any sensible inference. Our FRAME model, on the contrary, favors the specification of the joint distribution and characterizes local interactions by introducing non-linear functions of filter responses. This is not restricted by the neighborhood size since every filter introduces the same number of parameters regardless of its size, which enables us to explore structures at large scales (e.g., 33 \Theta 33 for the fabric texture). Moreover, FRAME can easily incorporate local interactions at different scales and orientations. It is also helpful to appreciate the difference between FRAME and the Gibbs distribution although both focus on the joint distributions. The Gibbs distribution is specified via potentials of various cliques, and the fact that most physical systems only have pair potentials (i.e., no potentials from the cliques with more than two pixels) is another reason why most MRF models for textures are restricted to auto-models. FRAME , on the other hand, builds potentials from finite-support filters and emphasizes the marginal distributions of filter responses. Although it may take a large number of filters to model a wide variety of textures, when it comes to modeling a certain texture, only a parsimonious set of the most meaningful filters needs to be selected. This selectivity greatly reduces the parameter size, thus allows accurate inference and modest computing. So FRAME is like a language: it has an efficient vocabulary (of filters) capable of describing most entities (textures), and when it comes to a specific entity, a few of the most meaningful words (filters) can be selected from the vocabulary for description. This is similar to the visual coding theory (Barlow et al 1989, Field 1989) which suggests that the sparse coding scheme has advantages over the compact coding scheme. The former assumes non-Gaussian distributions for f(I), whereas the latter assumes Gaussian distributions. Compared to the filtering method, FRAME has the following advantages: 1) solid statistical modeling, 2) it does not rely on the reversibility or reconstruction of I from fI (ff) g, and thus the filters can be designed freely. For example, we can use both linear and nonlinear filters, and the filters can be highly correlated to each other, whereas in the filtering method, a major concern is whether the filters form a tight frame (Daubechies 1992). There are various classifications for textures with respect to various attributes, such as Fourier and non-Fourier corresponding to whether the textures show periodic appearance; deterministic and stochastic corresponding to whether the textures can be characterized by some primitives and placement rules; and macro- and micro-textures in relation to the scales of local structures. FRAME erases these artificial boundaries and characterizes them in a unified model with different filters and parameter values. It has been well recognized that the traditional MRF models, as special cases of FRAME, can be used to model stochastic, non-Fourier micro-textures. From the textures we synthesized, it is evident that FRAME is also capable of modeling periodic and deterministic textures (fabric and pulses), textures with large scale elements (fur and cheetah blob), and textures with distinguishable textons (circles and cross bars), thus it realizes the full potential of MRF models. But the FRAME model is computationally very expensive. The computational complexity of the FRAME model comes from two major aspects. I). When bigger filters are adopted to characterize low resolution features, the computational cost will increase proportionally with the size of the filter window. II). The marginal distributions are estimated from sampled images, which requires long iterations for high accuracy of estimation. One promising way to reduce the computational cost is to combine the pyramid representation with the pseudo-likelihood estimation (Besag 1977). The former cuts the size of low resolution filters by putting them at the high levels of the pyramid as did in (Popat and Picard 1993), and the latter approximates E p [H (ff) ] by pseudo-likelihood and thus avoid the sampling process. But this method shall not be studied in this paper. No doubt many textures will not be easy to model, for example some human synthesized textures, such as textures on oriental rugs and clothes. It seems that the synthesis of such textures requires far more sophisticated or high-level features than those we used in this paper, and these high-level features may correspond to high-level visual process. At the same time, many theoretical issues remain yet to be fully understood, for example, the convergence properties of the sampling process and the definition of the best sampling procedures; the relationship between the sampling process and the physical process which forms the textures of nature and so on; and how to apply this texture model to the image segmentation problem (Zhu and Yuille 1996). It is our hope that this work will simulate future research efforts in this direction. Appendix . Filter pursuit and minimax entropy. This appendix briefly demonstrates the relationship between the filter pursuit method and the principle (Zhu, Wu and Mumford 1996). Let p(I; K ; SK ) be the maximum entropy distribution obtained at step k (see equation (18)), since our goal is to estimate the underlying distribution f(I), the goodness of p(I; K ; SK ) can be measured by the Kullback-Leibler distance between p(I; K ; SK ) and f(I) (Kullback and Leibler Z f(I) log f(I) it can be shown that As entropy(f(I)) is fixed, to minimize KL(f; p(I; K ; SK )) we need to choose SK such that has the minimum entropy, while given the selected filter set SK , p(I; K ; SK ) is computed by maximizing entropy(p(I)). In other words, for a fixed filter number K, the best set of filters is chosen by p2\Omega K where\Omega K is defined as equation (14). We call equation (29) the minimax entropy principle (Zhu, Wu, Mumford 1996). A stepwise greedy algorithm to minimize the entropy proceeds as the following. At step k suppose we choose F (fi) , and obtain the ME distribution p(I; f (ff) for fi. Then the goodness of F (fi) is measured by the decrease of the Kullback-Leibler distance KL(f(I); p(I; k It can be shown that where M is a covariance matrix of H (fi) , for details see (Zhu,Wu, Mumford 1996). Equation (31) measures a distance between f (fi) and E p(I; k;Sk ) [H (fi) ] in terms of variance, and therefore suggests a new form for the distance D(E p(I; k ;Sk equation (26), and this new form emphasizes the tails of the marginal distribution where important texture features lies, but the computational complexity is higher than the L 1 -norm distance. So far we have shown the filter selection in algorithm 3 is closely related to a minimax entropy principle. Acknowledgments This work was supported by the NSF grant DMS-91-21266 to David Mumford. The second author was supported by a grant to D.B. Rubin. --R "Finding minimum entropy codes." "Theories of visual texture perception." "Spatial interaction and the statistical analysis of lattice systems (with discussion)." "Efficiency of pseudolikelihood estimation for simple Gaussian fields." "Orthogonal distribution analysis: a new approach to the study of texture perception." "Entropy based algorithms for best basis selection." "Markov random field texture models." Ten lectures on wavelets. "On the Statistics of Vision: the Julesz Conjecture" "Ideal de-noising in an orthonormal basis chosen from a libary of bases. " "Relations between the statistics of natural images and the response properties of cortical cells" "Theory of communication." "Stochastic relaxation, Gibbs distribution, and the Bayesian restoration of images." "Markov random field image models and their applications to computer vision." "Annealing Markov chain Monto Carlo with applications to ancestral inference." "Statistics and structural approach to texture." "Pyramid-based texture analysis/synthesis." "Unsupervised texture segmentation using Gabor filters." "Information theory and statistical mechanics" "Visual pattern discrimination." "On information and sufficiency" Image representation using 2D Gabor wavelets. "Mumtiresolution approximations and wavelet orthonormal bases of L 2 (R)." Texture classification and segmentation using multiresolution simultaneous autoregressive models. "Time series models for texture synthesis." "Novel cluster-based probability model for texture synthesis, classification, and compression." "Multidimensional Markov chain models for image tex- tures." "Spatial-frequency organization in primate striate cortex." "Object and texture classification using higher order statistics." Image Analysis "Reaction-diffusion textures." "Spectral estimation for random fields with applications to Markov modeling and texture Classification." "Region Competition: unifying snakes, region growing, and Bayes/MDL for multi-band image segmentation" "Minimax Entropy Principle and its applications" --TR --CTR Dmitri Bitouk , Michael I. 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texture analysis and synthesis;minimax entropy;texture modeling;markov random field;feature pursuit;visual learning;maximum entropy

290114

A Router Architecture for Real-Time Communication in Multicomputer Networks.

AbstractParallel machines have the potential to satisfy the large computational demands of real-time applications. These applications require a predictable communication network, where time-constrained traffic requires bounds on throughput and latency, while good average performance suffices for best-effort packets. This paper presents a new router architecture that tailors low-level routing, switching, arbitration, flow-control, and deadlock-avoidance policies to the conflicting demands of each traffic class. The router implements bandwidth regulation and deadline-based scheduling, with packet switching and table-driven multicast routing, to bound end-to-end delay and buffer requirements for time-constrained traffic while allowing best-effort traffic to capitalize on the low-latency routing and switching schemes common in modern parallel machines. To limit the cost of servicing time-constrained traffic, the router includes a novel packet scheduler that shares link-scheduling logic across the multiple output ports, while masking the effects of clock rollover on the represention of packet eligibility times and deadlines. Using the Verilog hardware description language and the Epoch silicon compiler, we demonstrate that the router design meets the performance goals of both traffic classes in a single-chip solution. Verilog simulation experiments on a detailed timing model of the chip show how the implementation and performance properties of the packet scheduler scale over a range of architectural parameters.

Introduction Real-time applications, such as avionics, industrial process control, and automated manufacturing, impose strict timing requirements on the underlying computing system. As these applications grow in size and complexity, parallel processing plays an important role in satisfying the large computational demands. Real-time parallel computing hinges on effective policies for placing and scheduling communicating tasks in the system to ensure that critical operations complete by their deadlines. Ultimately, a parallel or distributed real-time system relies on an inter-connection network that can provide throughput and delay guarantees for critical communication between cooperating tasks; this communication may have diverse performance requirements, depending on the application [1]. However, instead of guaranteeing bounds on worst-case communication latency, most existing multicomputer network designs focus on providing good average network throughput and The work reported in this paper was supported in part by the National Science Foundation under grant MIP-9203895 and the Office of Naval Research under grants N00014-94-1-0229. Any opinions, findings, and conclusions or recommendations expressed in this paper are those of the authors and do not necessarily reflect the views of NSF or ONR. J. Rexford is with AT&T Labs - Research in Florham Park, New Jersey, and J. Hall and K. G. Shin are with the University of Michigan in Ann Arbor, Michigan. packet delay. Consequently, recent years have seen increasing interest in developing interconnection networks that provide performance guarantees in parallel machines [2-8]. Real-time systems employ a variety of network architec- tures, depending on the application domain and the performance requirements. Although prioritized bus and ring networks are commonly used in small-scale real-time systems [9], larger applications can benefit from the higher bandwidth available in multi-hop topologies. In addition, multi-hop networks often have several disjoint routes between each pair of processing nodes, improving the ap- plication's resilience to link and node failures. However, these networks complicate the effort to guarantee end-to- end performance, since the system must bound delay at each link in a packet's route. To deliver predictable communication performance in multi-hop networks, we present a novel router architecture that supports end-to-end delay and throughput guarantees by scheduling packets at each network link. Our prototype implementation is geared toward two-dimensional meshes, as shown in Figure 1; such topologies have been widely used as the interconnection network for a variety of commercial parallel machines. The design directly extends to a broad set of topologies, including the class of k-ary n-cube networks; with some changes in the routing of best-effort traffic, the proposed architecture applies to arbitrary point-to-point topologies. Communication predictability can be improved by assigning priority to time-constrained traffic or to packets that have experienced large delays earlier in their routes [10]. Ultimately, though, bounding worst-case communication latency requires prior reservation of link and buffer resources, based on the application's anticipated traffic load. Under this traffic contract, the network can provide end-to-end performance guarantees through effective link-scheduling and buffer-allocation policies. To handle a wide range of bandwidth and delay requirements, the real-time router implements the real-time channel [11- 13] abstraction for packet scheduling, as described in Section II. Conceptually, a real-time channel is a unidirectional virtual connection between two processing nodes, with a source traffic specification and an end-to-end delay bound. Separate parameters for bandwidth and delay permit the model to accommodate a wider range and larger number of connections than other service disciplines [14- 16], at the expense of increased implementation complexity. The real-time channel model guarantees end-to-end performance through a combination of bandwidth regulation and deadline-based scheduling at each link. Implementing packet scheduling in software would impose a significant burden on the processing resources at each node and Router to/from processor Fig. 1. Router in a Mesh Network: This figure shows a router in a 4 \Theta 4 square mesh of processing nodes. To communicate with another node, a processor injects a packet into its router; then, the packet traverses one or more links before reaching the reception port of the router at the destination node. would prove too slow to serve multiple high-speed links. This software would have to rank packets by deadline for each outgoing link, in addition to scheduling and executing application tasks. With high-speed links and tight timing constraints, real-time parallel machines require hardware support for communication scheduling. An efficient, low-cost solution requires a design that integrates this run-time scheduling with packet transmission. Hence, we present a chip-level router design that handles bandwidth regulation and deadline-based scheduling, while relegating non-realtime operations (such as admission control and route se- lection) to the network protocol software. Although deadline-based scheduling bounds the worst-case latency for time-constrained traffic, real-time applications also include best-effort packets that do not have stringent performance requirements [10, 11, 15, 17]; for ex- ample, good average delay may suffice for some status and monitoring information, as well as the protocol for establishing real-time channels. Best-effort traffic should be able to capitalize on the low-latency communication techniques available in modern parallel machines without jeopardizing the performance guarantees of time-constrained pack- ets. Section III describes how our design tailors network routing, switching, arbitration, flow-control, and deadlock- avoidance policies to the conflicting performance requirements of these two traffic classes. Time-constrained traffic employs packet switching and small, fixed-sized packets to bound worst-case performance, while best-effort packets employ wormhole switching [18] to reduce average latency and minimize buffer space requirements, even for large packets. The router implements deadlock-free, dimension- ordered routing for best-effort packets, while permitting the protocol software to select arbitrary multicast routes for the time-constrained traffic; together, flexible routing and multicast packet forwarding provide efficient group communication between cooperating real-time tasks. Section IV describes how the network can reserve buffer and link resources in establishing time-constrained connec- tions. In addition to managing the packet memory and connection data structures, the real-time router effectively handles the effects of clock rollover in computing scheduling for each packet. The router overlaps communication scheduling with packet transmission to maximize utilization of the network links. To reduce hardware complex- ity, the architecture shares packet buffers and sorting logic amongst the router's multiple output links, as discussed in Section V; a hybrid of serial and parallel comparison operations enables the scheduler to trade space for time to further reduce implementation complexity. Section VI describes the router implementation, using the Verilog hardware description language and the Epoch silicon compiler. The Epoch implementation demonstrates that the router can satisfy the performance goals of both traffic classes in an affordable, single-chip solution. Verilog simulation experiments on a detailed timing model of the chip show the correctness of the design and investigate the scaling properties of the packet scheduler across a range of architectural parameters. Section VII discusses related work on real-time multicomputer networks, while Section VIII concludes the paper with a summary of the research contributions and future directions. II. Real-Time Channels Real-time communication requires advance reservation of bandwidth and buffer resources, coupled with run-time scheduling at the network links. The real-time channel model [11] provides a useful abstraction for bounding end- to-end network delay, under certain application traffic characteristics Traffic parameters: A real-time channel is a unidirectional virtual connection that traverses one or more net-work links. In most real-time systems, application tasks exchange messages on a periodic, or nearly periodic, ba- sis. As a result, the real-time channel model characterizes each connection by its minimum spacing between messages resulting in a maximum transfer rate of Smax =I min bytes per unit time. To permit some variation from purely periodic traffic, a connection can generate a burst of up to Bmax messages in excess of the periodic restriction I min . Together, these three parameters form a linear bounded arrival process [19] that governs a connection's traffic generation at the source node. End-to-end delay bound: In addition to these traffic pa- rameters, a connection has a bound D on end-to-end message delay, based on the minimum message spacing I min . At the source node, a message m i generated at time t i has a logical arrival time ae By basing performance guarantees on these logical arrival times, the real-time channel model limits the influence an ill-behaving or malicious connection can have on other traffic in the network. The run-time link scheduler guarantees that message m i reaches its destination node by its deadline Per-hop delay bounds: The network does not admit a new connection unless it can reserve sufficient buffer and bandwidth resources without violating the requirements of Traffic Data Structure Queue 1 On-time time-constrained traffic Priority queue (by deadline '(m)+d) Queue 2 Best-effort traffic First-in-first-out queue Queue 3 Early time-constrained traffic Priority queue (by logical arrival time '(m)) I Real-Time Channel Scheduling Model: Under the real-time channel model, each link transmits traffic from three scheduling queues. To provide delay guarantees to time-constrained connections, the link gives priority to the on-time time-constrained messages in Queue 1 over the best-effort traffic in Queue 2. Queue 3 serves as a staging area for holding any early time-constrained messages. existing connections [11, 20]. A connection establishment procedure decomposes the connection's end-to-end delay bound D into local delay bounds d j for each hop in its route such that d j - I min and D. Based on the local delay bounds, a message m i has a logical arrival time at node j in its route, where j =0 corresponds to the source node. Link scheduling ensures that message m i arrives at node j no later than time local deadline at node j \Gamma1. In fact, message m i may reach node earlier, due to variations in delay at previous hops in the route. The scheduler at node j ensures that such "early" arrivals do not interfere with the transmission of "on-time" messages from other connections. Run-time link scheduling: Each link schedules time-constrained traffic, based on logical arrival times and dead- lines, in order to bound message delay without exceeding the reserved buffer space at intermediate nodes. The sched- uler, which employs a multi-class variation of the earliest due-date algorithm [21], gives highest priority to time-constrained messages that have reached their logical arrival time (i.e., ' j transmitting the message with the smallest deadline ' j , as shown in Table I. If Queue 1 is empty, the link services best-effort traffic from Queue 2, ahead of any early time-constrained messages This improves the average performance of best-effort traffic without violating the delay requirements of time-constrained communication. Queue 3 holds early time-constrained traffic, effectively absorbing variations in delay at the previous node. Upon reaching its logical arrival time, a message moves from Queue 3 to Queue 1. Link horizon parameter: By delaying the transmission of early time-constrained messages, the link scheduler can avoid overloading the buffer space at the downstream node [11, 15, 16]. Still, the scheduler could potentially improve link utilization and average latency by transmitting early messages from Queue 3 when the other two scheduling queues are empty. To balance this trade-off between buffer requirements and average performance, the link can transmit an early time-constrained message from Queue 3, as long as the message is within a small horizon h - 0 of its logical arrival time (i.e., ' j values of h permit the link to transmit more early time-constrained traffic, at the expense of increased memory requirements at the downstream node. Although each connection could conceivably have its own h value, employing a single horizon parameter allows the link to transmit early traffic directly from the head of Queue 3, without any per-connection data structures. Buffer requirements: To avoid buffer overflow or message loss, a connection must reserve sufficient memory for storing traffic at each node in its route. The required buffer space at node j depends on the connection's local delay bound d j , as well as the horizon parameter h j \Gamma1 for the incoming link. In particular, node j can receive a message from node j \Gamma 1 as early as ' j transmits the message at the earliest possible time. In the worst case, node j can hold a message until its deadline ' j . Hence, for this connection, messages m i stored at node j at time t. If a connection has messages arrive as early as possible, and depart as late as possible, then node could have to store as many as I min messages from this connection at the same time. By reserving buffer and bandwidth resources in advance, the real-time channel model guarantees that every message arrives at its destination node by its deadline, independent of other best-effort and time-constrained traffic in the network. III. Mixing Best-Effort and Time-Constrained Traffic Although the real-time channel model bounds the worst-case performance of time-constrained messages, the scheduling model in Table I can impose undue restrictions on the packet size and flow-control schemes for best-effort traffic. To overcome these limitations, we propose a router architecture that tailors its low-level communication policies to the unique demands of the two traffic classes. Fine- grain, priority-based arbitration at the network links permits the best-effort traffic to capitalize on the low-latency techniques in modern multicomputer networks without sacrificing the performance guarantees of the time-constrained connections. Figure 2 shows the high-level architecture of the real-time router, with separate control and data path for the two traffic classes. incoming links Address reception injection control Packet Scheduling Logic new address Memory Packet outgoing links Free Table Connection best-effort data connection id address time-constrained best-effort time-constrained Fig. 2. Real-Time Router: This figure shows the real-time router architecture, with separate control and data path for best-effort and time-constrained packets. The router includes a packet memory, connection routing table, and scheduling logic to support delay and bandwidth guarantees for time-constrained traffic. To connect to the local processor, the router exports a control interface, a reception port, and separate injection ports for each traffic class. A. Complementary Switching Schemes To ensure that time-constrained connections meet their delay requirements, the router must have control over bandwidth and memory allocation. For example, suppose that a time-constrained message arrives with a tight deadline (i.e., '(m i while the outgoing link is busy transmitting other traffic. To satisfy this tight timing requirement, the outgoing link must stop servicing any lower-priority messages within a small, bounded amount of time. This introduces a direct relationship between connection admissibility and the maximum packet size of the time-constrained and best-effort traffic sharing the link. In most real-time systems, time-constrained communication consists of 10-20 byte exchanges of command or status information [9]. Consequently, the real-time router restricts time-constrained traffic to small, fixed-size packets that can support a distributed memory read or write operation. This bounds link access latency and buffering delay while simplifying memory allocation in the router. To ensure predictable consumption of link and buffer re- sources, time-constrained traffic employs store-and-forward packet switching . By buffering packets at each node, packet switching allows each router to independently schedule packet transmissions to satisfy per-hop delay requirements. To improve average performance, the time-constrained traffic could conceivably employ virtual cut-through switching [22] to allow an incoming packet to proceed directly to an idle outgoing link. However, in contrast to traditional virtual cut-through switching of best effort traf- fic, the real-time router cannot forward a time-constrained packet without first assessing its logical arrival time (to ensure that the downstream router has sufficient buffer space for the packet) and computing the packet deadline (which serves as the logical arrival time at the down-stream router). To avoid this extra complexity and over- head, the initial design of the real-time router implements store-and-forward packet switching, which has the same worst-case performance guarantees as virtual cut-through switching. A future implementation could employ virtual cut-through switching to reduce the average latency of the time-constrained traffic. Although packet switching delivers good, predictable performance to small, time-constrained packets, this approach would significantly degrade the average latency of long, best-effort packets. Even in a lightly-loaded network, end-to-end latency under packet switching is proportional to the product of packet size and the length of the route. Instead, the best-effort traffic can employ wormhole switching [18] for lower latency and reduced buffer space require- ments. Similar to virtual cut-through switching, wormhole switching permits an arriving packet to proceed directly to the next node in its route. However, when the outgoing link is not available, the packet stalls in the network instead of buffering entirely within the router. In effect, wormhole switching converts the best-effort scheduling "queue" in Table I into a logical queue that spans multiple nodes. The router simply includes small five-byte flit (flow control unit) buffers [23] to hold a few bytes of a packet from each input link. When an incoming packet fills these buffers, inter-node flow control halts further transmission from the previous node until more space is available; once the five-byte chunk proceeds to a buffer at the outgoing link, the router transmits an acknowledg- 1virtual channel id data byte strobe/enable flit acknowledgement Fig. 3. Link Encoding: In the real-time router, each link can transmit a byte of data, along with a strobe signal and a virtual channel identifier. In the reverse direction, an acknowledgment bit indicates that the router can store another flit on the best-effort virtual channel. x offset y offset length data bytes connection id data bytes (18) (a) Best-effort packet (b) Time-constrained packet Fig. 4. Packet Formats: This figure illustrates the packet formats for best-effort and time-constrained packets in the real-time router. Best-effort packets consist of a two-byte routing header and a one-byte length field, along with the variable-length data. Time-constrained packets are 20 bytes long and include the connection identifier and the deadline from the previous hop in the route, which serves as the logical arrival time at the current router. ment bit to signal the upstream router to start sending the next flit. This fine-grain, per-hop flow control permits best-effort traffic to use large variable-sized packets, reducing or even avoiding packetization overheads, without increasing buffer complexity in the router. The combination of wormhole and packet switching, with best-effort traffic consuming small flit buffers and time-constrained connections reserving packet buffers, results in an effective partitioning of router resources. B. Separate Logical Resources Even though wormhole and packet switching exercise complementary buffer resources, best-effort and time-constrained traffic still share access to the same net-work links. To provide tight delay guarantees for time-constrained connections, the router must bound the time that the variable-sized, wormhole packets can stall the forward progress of on-time, time-constrained traffic. How- ever, a blocked wormhole packet can hold link resources at a chain of consecutive routers in the network, indirectly delaying the advancement of other traffic that does not even use the same links. This complicates the effort to provision the network to bound worst-case end-to-end latency, as discussed in the treatment of related work in Section VII. In order to control the interaction between the two traffic classes, the real-time router divides each link into two virtual channels [23]. A single bit on each link differentiates between time-constrained and best-effort packets, as shown in Figure 3; each link also includes an acknowledgment bit for flow control on the best-effort virtual channel. Each wormhole virtual channel performs round-robin arbitration on the input links to select an incoming best-effort packet for service, while the packet-switched virtual channel transmits time-constrained packets based on their deadlines and logical arrival times. Priority arbitration between the two virtual channels tightly regulates the intrusion of best-effort traffic on time-constrained packets on each outgoing link. This effectively provides flit-level pre-emption of best-effort traffic whenever an on-time time-constrained packet awaits service, while permitting wormhole flits to consume any excess link bandwidth. In a separate simulation study, we have demonstrated the effectiveness of using flit-level priority arbitration policies to mix best-effort wormhole traffic and time-constrained packet-switched traffic [24-26]. While the real-time router gives preferential treatment to time-constrained traffic, the outgoing links transmit best-effort flits ahead of any early time-constrained packets, consistent with the policies in Table I. Although this arbitration mechanism ensures effective scheduling of the traffic on the outgoing links and the reception port, the best-effort and time-constrained packets could still contend for resources at the injection port at the source node. The local processor could solve this problem by negotiating between best-effort and time-constrained traffic at the injection port, but this would require the processor to perform flit-level arbitration. Instead, the real-time router includes a dedicated injection port for each traffic class. The two injection ports, coupled with the low-level arbitration on the outgoing links, ensure that time-constrained traffic has fine-grain preemption over the best-effort packets across the entire path through the network, while allowing best-effort packets to capitalize on any remaining link band-width C. Buffering and Packet Forwarding To support the multiple incoming and outgoing ports, the real-time router design requires high throughput for re- ceiving, storing, and transmitting packets. Internally, the router isolates the best-effort and time-constrained traffic on separate buses to increase the throughput and reduce the complexity of the arbitration logic. Each incoming and outgoing port includes nominal buffer space to avoid stalling the flow of data while waiting for access to the bus. The best-effort bus is one flit wide and performs round-robin arbitration among the flit buffers at the incoming ports. Running at the same speed as the byte-wide input ports, this five-byte bus has sufficient throughput to accommodate a peak load of best-effort traffic. Transferring best-effort packets in five-byte chunks incurs a small initial transmission delay at each router, which could be reduced by using a crossbar switch; however, we employ a shared bus for the sake of simplicity. Other recent multi-computer router architectures have used a wide bus for flit transfer [27, 28]. The structure and placement of packet buffers plays a large role in the router's ability to accommodate the performance requirements of time-constrained connections. The simplest solution places a separate queue at each input link. However, input queuing has throughput limitations [29], since a packet may have to wait behind other traffic destined for a different outgoing link. In addition, queuing packets at the incoming links complicates the effort to schedule outgoing traffic based on delay and throughput requirements. Instead, the real-time router queues time-constrained packets at the output ports; the router shares a single packet memory among the multiple output ports to maximize the network's ability to accommodate time-constrained connections with diverse buffer requirements. To accommodate the aggregate memory bandwidth of the five input and five output ports, the router stores packets in 10-byte chunks, with demand-driven round-robin arbitration amongst the ports. Since time-constrained traffic is not served in a first-in first-out order, the real-time router must have a data structure that records the idle memory locations in the packet buffer. Similar to many shared-memory switches in high-speed networks, the real-time router maintains an idle-address pool [29], implemented as a stack. This stack consists of a small memory, which stores the address of each free location in the packet buffer, and a pointer to the first entry. Initially, the stack includes the address of each location in the packet memory. An incoming packet retrieves an address from the top of the stack and increments the stack pointer to point to the next available entry. Upon packet departure, the router decrements this pointer and returns the free location to the top of stack. The idle- address stack always has at least one free address when a new packet arrives, since the real-time channel model never permits the time-constrained traffic to overallocate the buffer resources. D. Routing and Deadlock-Avoidance Although wormhole switching reduces the buffer requirements and average latency for best-effort traffic, the low-level inter-node flow control could potentially introduce cyclic dependencies between stalled best-effort packets. To avoid these cycles, the real-time router implements dimension-ordered routing, a shortest-path scheme that completely routes a packet in the x-direction before proceeding in the y-direction to the destination, as shown by the shaded nodes in Figure 1. Dimension-ordered routing avoids packet deadlock in a square mesh [30] and also facilitates an efficient implementation based on x and y offsets in the packet header, as shown in Figure 4(a); the offsets reach zero when the packet has arrived at its destination node. To improve the performance of best-effort traffic, an enhanced version of the router could support adaptive wormhole routing and additional virtual channels, at the expense of increased implementation complexity [31, 32]. In particular, non-minimal adaptive routing would enable best-effort packets to circumvent links with a heavy load of time-constrained traffic. Although routing is closely tied with deadlock-avoidance for best-effort packets, the real-time router need not dictate a particular routing scheme for the time-constrained traf- fic. Instead, each time-constrained connection has a fixed path through the network, based on a table in each router; this table is indexed by the connection identifier field in the header of each time-constrained packet, as shown in Figure 4(b). As part of establishing a real-time channel, the network protocol software can select a fixed path from the source to the destination(s), based on the available band-width and buffer resources at the routers. The protocol software can employ a variety of algorithms for selecting unicast and multicast routes based on the resources available in the network [33]. Once the connection establishment protocol reserves buffer and bandwidth resources for a real-time channel, the combination of bandwidth regulation and packet scheduling prevents packet deadlock for time-constrained traffic. Table II summarizes how the real-time router employs these and other policies to accommodate the conflicting performance requirements of the two traffic classes. IV. Managing Time-Constrained Connections A real-time multicomputer network must have effective mechanisms for establishing connections and scheduling packets, based on the delay and throughput requirements of the time-constrained traffic. To permit a single-chip im- plementation, the real-time router offloads non-real-time operations, such as route selection and admission control, to the network protocol software. At run-time, the router coordinates access to buffer and link resources by managing the packet memory and the connection data struc- tures. In addition, the router architecture introduces efficient techniques for bounding the range of logical arrival times and deadlines, to limit scheduler delay and implementation complexity. A. Route Selection and Admission Control Establishing a real-time channel requires the application to specify the traffic parameters and performance requirements for the new connection. Admitting a new con- nection, and selecting a multi-hop route with suitable local delay parameters, is a computationally-intensive procedure [10, 11, 20]. Fortunately, channel establishment typically does not impose tight timing constraints, in contrast to the actual data transfer which requires explicit guarantees on minimumthroughput and worst-case delay. In fact, in most cases, the network can establish the required time-constrained connections before the application commences. To permit a single-chip solution, the real-time router relegates these non-real-time operations to the protocol soft- ware. The network could select routes and admit new connections through a centralized server or a distributed pro- tocol. In either case, this protocol software can use the best-effort virtual network, or even a set of dedicated time-constrained connections, to exchange information to select a route and provision resources for each new connection. The route selected for a connection depends on the traffic characteristics and performance requirements, as well as the available buffer and bandwidth resources in the net- work. As part of establishing a new real-time channel, the protocol software assigns a unique connection identifier at each hop in the route. Then, each node in the route writes Time-Constrained Best-Effort Switching Packet switching Wormhole switching Packet size Small, fixed size Variable length Link arbitration Deadline-driven Round-robin on input links Routing Table-driven multicast Dimension-ordered unicast Buffers Shared output queues Flit buffers at input links Flow control Rate-based Flit acknowledgments II Architectural Parameters: This table summarizes how the real-time router supports the conflicting performance requirements of time-constrained and best-effort traffic. Write Command Fields Connection parameters outgoing connection id local delay bound d bit-mask of output ports incoming connection id Horizon parameter bit-mask of output ports horizon value h III Control Interface Commands: This table summarizes the control commands used to configure the real-time router. control information into the router's connection table, as shown in Table III. At run-time, this table is indexed by the connection identifier field of each incoming time-constrained packet, as shown in Figure 4(b). To minimize the number of pins on the router chip, the controlling processor updates this table as a sequence of four, one-byte operations that specify the incoming connection identifier and the three fields in the table. After closing a connection, the network protocol software can reuse the connection identifier by overwriting the entry in the routing table. The processor uses the same control interface to set the horizon parameters h for each of the five outgoing ports. As shown in Table III, the routing table stores the con- nection's identifier at the next node, the local delay bound d, and a bit mask for directing traffic to the appropriate outgoing port(s). When a packet arrives, the router indexes the table with the incoming connection identifier and replaces the header field with the new identifier for the down-stream router. At the same time, the router computes the packet's deadline from the logical arrival time in the packet header and the local delay bound in the connection ta- ble. Finally, the bit mask permits the router to forward an incoming packet to multiple outgoing ports, allowing the network protocol software to establish multicast real-time channels. This facilitates efficient, timely communication between a set of cooperating nodes. To simplify the design, the real-time router requires a multicast connection to use the same value of d for each of its outgoing ports at a single node. Then, based on the bit mask in the routing table, the router queues the updated packet for transmission on the appropriate outgoing port(s). By implementing a shared packet memory, the real-time router can store a single copy of each multicast packet, removing the packet only after it has been transmitted by each output port selected in the bit mask. The shared packet memory also permits the network protocol software to employ a wide variety of buffer allocation policies. On the one extreme, the route selection and admission control protocols could allocate packet buffers to any new connec- tion, independent of its outgoing link. However, this could allow a single link to consume the bulk of the memory loca- tions, reducing the chance of establishing time-constrained connections on the other outgoing links. Instead, the admission control protocol should bound the amount of buffer space available to each of the five outgoing ports. Simi- larly, the network could limit the size of the link horizon parameters h to reduce the amount of memory required by each connection. In particular, at run-time, a higher-level protocol could reduce the h values of a router's incoming links when the node does not have sufficient buffer space to admit new connections. B. Handling a Clock with Finite Range The packet deadline at one node serves as the logical arrival time at the downstream node in the route. Carrying these logical arrival times in the packet header, as shown in Figure 4(b), implicitly assumes that the net-work routers have a common notion of time, within some bounded clock skew. Although this is not appropriate in a wide-area network context, the tight coupling in parallel machines minimizes the effects of clock skew. Alternatively, the router could store additional information in the connection table to compute ' j ) from a packet's actual arrival time and the logical arrival time of the connection's previous packet [34]; however, this approach would require the router to periodically refresh this connection state to correctly handle the effects of clock rollover. Instead, the real-time router avoids this overhead by capitalizing on the tight coupling between nodes to assume synchronized clocks. Even with synchronized clocks, the real-time router cannot completely ignore the effects of clock rollover. To schedule time-constrained traffic, the router architecture includes a real-time clock, implemented as a counter that increments once per packet transmission time. For a practical implementation, the router must limit the number of bits b used to represent the logical arrival times and deadlines of time-constrained packets. Since logical arrival times continually increase, the design must use modulo arithmetic to compute packet deadlines and schedule traffic for transmission. As a result, the network must restrict the logical arrival times that can exist in a router at the same time; otherwise, the router cannot correctly distinguish between different packets awaiting access to the outgoing link. Selecting a value for b introduces a fundamental trade-off between connection admissibility and scheduler complex- ity. To select a packet for transmission, the scheduler must compare the deadlines and logical arrival times of the time-constrained packets; for example, the data structures in Table I require comparison operations to enqueue/dequeue packets. Larger values of b would increase the hardware cost and latency for performing these packet comparison operations. However, smaller values of b would restrict the network's ability to select large delay bounds d and horizon parameters h for time-constrained connections. The network protocol software can limit the delay and horizon parameters, based on the value of b imposed by the router implementation. Alternatively, in implementing the router, a designer could select a value for b based on typical requirements for the expected real-time applications. To formalize the trade-off between complexity and ad- missibility, consider a connection traversing consecutive links local delay parameters d j \Gamma1 and d j respectively, where link As discussed in Section II, a packet can arrive as much as units ahead of its logical arrival time ' j and depart as late as its deadline ' j . Consequently, for any messages m i from this connection at time t. The network must ensure that the router can differentiate between the full range of logical arrival times in this set. The router can correctly interpret logical arrival times and deadlines, even in the presence of clock rollover, as long as every connection has values that are less than half the range of the on-chip clock. That is, the router requires d j connections sharing the link. Under this restriction, the router can compare packets based on their logical arrival times and deadlines by using modulo arithmetic. For example, suppose (i.e., the clock has a range of 256 time units) and the connections all configuration corresponds to Figure 5. Any early packets have logical arrival times between 240 and 346, modulo 256. For example, a packet with '(m)=80 would be considered early traffic (since Similarly, any on-time packets have logical arrival times between 200 and 240. For example, a packet with would be considered on-time traffic (since deadlines Hence, these deadlines also fall within the necessary range in Figure 5, allowing the router to compute (' j range of packet logical arrival times128 Fig. 5. Handling Clock Rollover: This figure illustrates the effects of clock rollover with an 8-bit clock, where the current time is t =240(mod 256). In the example, all connections satisfy d and d ensuring that the router can correctly compare to t to distinguish between on-time and early packets. compare on-time packets based on their deadlines. V. Scheduling Time-Constrained Packets To satisfy connection delay, throughput, and buffer requirements, each outgoing port must schedule time-constrained packets based on their logical arrival times and deadlines, as well as the horizon parameter. The real-time router reduces implementation complexity by sharing a single scheduler amongst the early and on-time traffic on each of the five output ports. Extensions to the scheduler architecture further reduce the implementation cost by trading space for time. A. Integrating Early and On-Time Packets To maximize link utilization and channel admissibility, each outgoing port should overlap packet scheduling operations with packet transmission. As a result, packet size determines the acceptable worst-case scheduling de- lay. Scheduling time-constrained traffic, based on delay or throughput parameters, typically requires a priority queue to rank the outgoing packets. Priority queue architectures introduce considerable hardware complexity [35-39], particularly when the link must handle a wide range of packet priorities or deadlines. For example, most high-speed solutions require O(n) hardware complexity to rank n packets, using a systolic array or shift register consisting of n comparators [35, 40, 41]. Additional technical challenges arise in trying to integrate packet scheduling with bandwidth regulation [42], since the link cannot transmit a packet unless it has reached its logical arrival time. To perform bandwidth regulation and deadline-based scheduling, the real-time router could include two priority queues for each of its five outgoing ports, as suggested by Table I. However, this approach would be extremely expensive and would require additional logic to transfer packets from the "early" queue to the "on-time" queue; this is particularly complicated when multiple packets reach their eligibility times simultaneously. In the worst case, an out-going port could have to dequeue a packet from Queue 1 or Queue 3, enqueue several arriving packets to Queue 1 adder logic (l> bit mask l l -t l adder logic on-time enable (from ineligible eligible/ check mask select port comparators horizon parameter Fig. 6. Comparator Tree Scheduler: This figure shows the scheduling architecture in the real-time router. The leaf nodes at the base of the comparator tree stores a small amount of per-packet state information. On-time: Fig. 7. Scheduler Keys: This figure illustrates how the real-time router assigns a key to each time-constrained packet awaiting transmission on an outgoing port. A single bit differentiates on-time and early packets; ineligible traffic refers to packets that are not destined to this port. and/or Queue 3, and move a large number of packets from Queue 3 to Queue 1, all during a single packet transmission time. To avoid this complexity, the real-time router does not attempt to store the time-constrained packets in sorted order. Instead, the router selects the packet with the smallest key via a comparator tree, as shown in Figure 6. Like the systolic and shift register approaches, the tree architecture introduces O(n) hardware complexity. For the moderate size of n in a single-chip router, the comparator tree can overlap the O(lgn) stages of delay with packet transmission. To avoid this excessive complexity, the real-time router integrates early and on-time packets into a single data structure. Each link schedules time-constrained packets based on sorting keys, as shown in Figure 7, where smaller have higher priority. A single bit differentiates between early and on-time packets. For on-time traffic, the lower bits of the key represent packet laxity , the time remaining till the local deadline expires, whereas the key for early traffic represents the time left before reaching the packet's logical arrival time. The packet keys are normal- ized, relative to current time t, to allow the scheduler to perform simple, unsigned comparison operations, even in the presence of clock rollover. Each scheduling operation operates independently to locate the packet with the minimum sorting key, permitting dynamic changes in the values of keys. The base of the tree computes a key for each packet, based on the packet state and the current time t, as shown in the right side of Figure 6; the base of the tree stores per-packet state information, whereas the packet memory stores the actual packet contents. B. Sharing the Scheduler Across Output Ports By using a comparator tree, instead of trying to store the packets in sorted order, the router can allow all five out-going ports to share access to this scheduling logic, since the tree itself does not store the packet keys. As shown in Figure 6, each leaf in the tree stores a logical arrival time '(m), a deadline '(m)+d, and a bit mask of outgoing ports, assigned at packet arrival based on the connection state. The bit mask determines if the leaf is eligible to compete for access to a particular outgoing port. When a port transmits a selected packet, it clears the corresponding field in the leaf's bit mask; a bit mask of zero indicates an empty packet leaf slot and a corresponding idle slot in the packet memory. The base of the tree also determines if packets are early ('(m) ? t) or on-time ('(m) - t) and computes the sorting keys based on the current value of t. At the top of the sorting tree, an additional comparator checks to see if the winner is an early packet that falls within the port's horizon parameter; if so, the port transmits this packet, unless best-effort flits await service. Still, to share the comparator logic, the scheduler must operate quickly enough to overlap run-time scheduling with packet transmission on each of the outgoing ports. Conse- quently, the real-time router pipelines access to the comparator tree. With p stages of pipelining, the scheduler has a row of latches at in the tree, to store the sorting key and buffer location for the winning packet in the subtrees. Every few cycles, another link begins its scheduling operation at the base of the tree. Similarly, every few horizon parameter levels packets Fig. 8. Logic Sharing: This figure illustrates how the scheduler can trade space for time by sharing comparator logic amongst groups of k packets. cycles, another link completes a scheduling operation and can initiate a packet transmission. As a result, the router staggers packet departures on the five outgoing ports. The necessary amount of pipelining depends on the latency of the comparator tree, relative to the packet transmission delay. C. Balancing Hardware Complexity and Scheduler Latency The pipelined comparator tree has relatively low hardware cost, compared to alternate approaches that implement separate priority queues for the early and on-time packets on each outgoing port. However, as shown in Section VI, the scheduler logic is still the main source of complexity in the real-time router architecture. To handle n packets, the scheduler in Figure 6 has a total of 2+ lg n stages of logic, including the operations at the base of the tree as well as the comparator for the horizon parameter. In terms of implementation cost, the tree requires n comparators and n leaf nodes, for a total of 2n elements of similar complexity. As n grows, the number of leaf nodes can have a significant influence on the bus loading at the base of the tree. Fortunately, for certain values of n, the comparator tree has low enough latency to avoid the need to fully pipeline the scheduling logic. This suggests that the scheduler could reduce the number of comparators by trading space for time. Under this approach, the scheduler combines several leaf units into a single module with a small memory (e.g., a register file) to store the deadlines and logical arrival times for k packets, as shown in Figure 8. At the base of the tree, each of the n=k modules can sequentially compare its k sorting keys, using a single comparator, to select the packet with the minimum incurs k stages of delay. Then, a smaller comparator tree finds the smallest key amongst n=k packets. As a result, the scheduler incurs stages of delay. Note that, for the architecture reduces to the comparator tree in Figure 6, with its 2 stages of logic. For larger values of k, the scheduler has larger arbitration delay but reduced implementation com- plexity. The architecture in Figure 8 has 2n=k compara- tors, as well as a lighter bus loading of n=k elements at the base of the tree. In addition, larger values of k allow the base of the tree to consist of n=k k-element register files, instead of n individual registers, with a reduction in chip complexity. With a careful selection of n and k, the real-time router can have an efficient, single-chip implementation that performs bandwidth regulation and deadline-based scheduling on multiple outgoing ports. VI. Performance Evaluation To demonstrate the feasibility of the real-time router, and study its scaling properties, a prototype chip has been designed using the Verilog hardware description language and the Epoch silicon compiler from Cascade Design Au- tomation. This framework facilitates a detailed evaluation of the implementation and performance properties of the architecture. The Epoch tools compile the structural and behavioral Verilog models to generate a chip layout and an annotated Verilog model for timing simulations. These tools permit extensive testing and performance evaluation without the expense of chip fabrication. A. Router Complexity Using a three-metal, 0:5-m CMOS process, the 123-pin chip has dimensions 8:1 mm \Theta 8:7 mm for an implementation with 256 time-constrained packets and up to 256 connections, as shown in Table IV. The scheduling logic accounts for the majority of the chip area, with the packet memory consuming much of the remaining space, as shown in Table V. Operating at 50 MHz, the chip can transmit or receive a byte of data on each of its ten ports every 20 nsec. This closely matches the access time of the 10-byte-wide, single-ported SRAM for storing time-constrained traffic; the memory access latency is the bottleneck in this realization of the router. Since time-constrained packets are 20-bytes long, the scheduling logic must select a packet for transmission every 400 nsec for each of the five output ports To match the memory and link throughputs, the comparator tree consists of a two-stage pipeline, where each stage requires approximately 50 nsec. Although the tree could incorporate up to five pipeline stages, the two-stage design provides sufficient throughput to satisfy the output ports. This suggests that the link scheduler could effectively support a larger number of packets or additional output ports, for a higher-dimensional mesh topology. Alternately, the router design could reduce the hardware cost of the comparator tree by sharing comparator logic between multiple leaves of the tree, as Parameter Value Connections 256 Time-constrained packets 256 (sorting Comparator tree pipeline 2 stages Flit input buffer 10 bytes (a) Architectural parameters Parameter Value Process 0.5-m 3-metal CMOS Signal pins 123 Transistors 905; 104 Area 8:1 mm \Theta 8:7 mm Power 2:3 watts (b) Chip complexity IV Router Specification: This table summarizes the architectural parameters and chip complexity of the prototype implementation of the real-time router. Unit Area Transistors Packet scheduler 34.02 mm 2 555025 Memory and control 5.97 mm 2 268161 Connection table 0.65 mm 2 20966 Idle-address pool 0.35 mm 2 15600 Router Components: This table summarizes the area contribution and transistor count for the main components of the router. discussed in Section V-C. Figure 9 highlights the cost-performance trade-offs of logic sharing, based on Epoch implementations and Verilog simulation experiments. As k increases, the scheduler complexity decreases in terms of area, transistor count, and power dissipation, with reasonable increases in scheduler latency. The results start with a grouping size of k= 4, since the Epoch library does not support static RAM components with fewer than four lines. 1, the graphs plot results from the router implementation in Table V, which uses flip-flops to store packet state at the base of the tree. The Epoch silicon compiler generates a better automated layout of these flip-flops than of the small SRAMs, resulting in better area statistics in Table V, despite the larger transistor count. A manual layout would significantly improve the area statistics for still, the area graph shows the relative improvement for larger values of k.) These plots can help guide the trade-off between hardware complexity and scheduler latency in the router imple- mentation. For example, a group size of k =4 reduces the number of transistors by 45% (from 555; 025 to 306; 829). The number of transistors does not decrease by a factor of four, since the smaller scheduler still has to store the state information for each packet; in addition, the scheduler requires additional logic and registers to serialize access to the shared comparators. Still, logic sharing significantly reduces implementation complexity. Larger values of k further reduce the number of comparators and improve the density of the memory at the base of the tree. Scheduler latency does not grow significantly for small values of k. For 4, delay in the comparator tree increases by just 67% (from 0:115 -sec to 0:192 -sec). The lower bus loading at the base of the tree helps counteract the increased latency from serializing access to the first layer of comparators and significantly reduces power dissipation. B. Simulation Experiments Since Verilog simulations of the full chip are extremely memory and CPU intensive, we focus on a modest set of timing experiments, aimed mainly at testing the correctness of the design. A preliminary experiment tests the baseline performance of best-effort wormhole packets. To study a multi-hop configuration, the router connects its links in the x and y directions. The packet proceeds from the injection port to the positive x link, then travels from the negative x input link to the positive y direction; after reentering the router on the negative y link, the packet proceeds to the reception port. In this test, a b byte wormhole packet incurs an end-to-end latency of where the link transmits one byte in each cycle. This delay is proportional to packet length, with a small overhead for synchronizing the arriving bytes, processing the packet header, and accumulating five-byte chunks for access to the router's internal bus. In contrast, packet switching would introduce additional delay to buffer the packet at each hop in its route. An additional experiment illustrates how the router schedules time-constrained packets to satisfy delay and throughput guarantees, while allowing best-effort traffic to capitalize on any excess link bandwidth. Figure 10 plots the link bandwidth consumed by best-effort traffic and each of three time-constrained connections with the following parameters, in units of 20-byte slots: d I min All three connections compete for access to a single network link with horizon parameter h= 0, where each connection has a continual backlog of traffic. The time-constrained connections receive service in proportion to their through-put requirements, since a packet is not eligible for service k (group size)100000300000500000Number of transistors Number of transistors Area (square millimeters) Area of tree (microseconds) Scheduling latency (milliwatts) Power dissipation Fig. 9. Evaluating Logic Sharing: These plots compare different implementations of the comparator tree, with different group sizes k. As k grows, implementation complexity decreases but scheduler latency increases. till its logical arrival time. Similarly, the link transmits each packet by its deadline, with best-effort flits consuming any remaining link bandwidth. VII. Related Work This paper complements recent work on support for real-time communication in parallel machines [2-7]. Several projects have proposed mechanisms to improve predictability in the wormhole-switched networks common in modern multicomputers. In the absence of hardware support for priority-based scheduling, application and operating system software can control end-to-end performance by regulating the rate of packet injection at each source node [7]. However, this approach must limit utilization of the communication network to account for possible contention between packets, even from lower-priority traffic. This is a particularly important issue in wormhole networks, since a stalled packet may indirectly block the advancement of other traffic that does not even use the same links. The underlying router architecture can improve predictability by favoring older packets when assigning virtual channels or arbitrating between channels on the same physical link [23]. Although these mechanisms reduce variability in end-to- Time (clock cycles)100300500 Connection service (bytes) best-effort connection 2 connection 1 connection 0 Fig. 10. Timing Experiment: This experiment evaluates a mixture of time-constrained and best-effort packets competing for access to a single outgoing link with horizon The scheduler satisfies the deadlines of the time-constrained packets, while permitting best-effort flits to capitalize on any additional band-width latency, more aggressive techniques are necessary to guarantee performance under high network utilization. A router can support multiple classes of traffic, such as user and system packets, by partitioning traffic onto different virtual channels, with priority-based arbitration for access to the network links [23]. Flit-level preemption of low-priority virtual channels can significantly reduce intrusion on the high-priority packets. Still, these coarse-grain priorities do not differentiate between packets with different latency tolerances. With additional virtual channels, the network has greater flexibility in assigning packet priority, perhaps based on the end-to-end delay requirement, and restricting access to virtual channels reserved for higher-priority traffic [4, 5]. Coupled with restrictions on the source injection rate, these policies can bound end-to-end packet latency by limiting the service and blocking times for higher-priority traffic [3]. Although assigning priorities to virtual channels provides some control over packet scheduling, this ties priority resolution to the number of virtual channels. The router can support fine-grain packet priorities by increasing the number of virtual channels, at the expense of additional implementation complexity; these virtual channels incur the cost of additional flit buffers and larger virtual channel identifiers, as well as more complex switching and arbitration logic [32]. Instead of dedicating virtual channels and flit buffers to each priority level, a router can increase priority resolution by adopting a packet-switched design. The priority-forwarding router chip [6] follows this approach by employing a 32-bit priority field in small, 8- packet priority queues at each input port. The router incorporates a priority-inheritance protocol to limit the effects of priority inversion when a full input buffer limits the transmission of high-priority packets from the previous node; the input buffer's head packet inherits the priority of the highest-priority packet still waiting at the up-stream router. In contrast, the real-time router implements a single, shared output buffer that holds up to 256 time-constrained packets, with a link-scheduling and memory reservation model that implicitly avoids buffer overflow. By dynamically assigning an 8-bit packet priority at each node, the real-time router can satisfy a diverse range of end-to- delay bounds, while permitting best-effort wormhole traffic to capitalize on any excess link bandwidth. VIII. Conclusion Parallel real-time applications impose diverse communication requirements on the underlying interconnection net- work. The real-time router design supports these emerging applications by bounding packet delay for time-constrained traffic, while ensuring good average performance for best-effort traffic. Low-level control over routing, switching, and flow control, coupled with fine-grain arbitration at the network links, enables the router to effectively mix these two diverse traffic classes. Careful handling of clock rollover enables the router to support connections with diverse delay and throughput parameters with small keys for logical arrival times and deadlines. Sharing scheduling logic and packet buffers amongst the five output ports permits a single-chip solution that handles up to 256 time-constrained packets simultaneously. Experiments with a detailed timing model of the router chip show that the design can operate at 50 MHz with appropriate pipelining of the scheduling logic. Further experiments show that the design can trade space for time to reduce the complexity of the packet scheduler. As ongoing research, we are considering alternate link- scheduling algorithms that would improve the router's scal- ability. In this context, we are investigating efficient hardware architectures for integrating bandwidth regulation and packet scheduling [42]; these algorithms include approximate scheduling schemes that balance the trade-off between accuracy and complexity, allowing the router to efficiently handle a larger number of time-constrained pack- ets. We are also exploring the use of the real-time router as a building block for constructing large, high-speed switches that support the quality-of-service requirements of real-time and multimedia applications. The router's delay and throughput guarantees for time-constrained traffic, combined with good best-effort performance and a single-chip implementation, can efficiently support a wide range of modern real-time applications, particularly in the context of tightly-coupled local area networks. --R "Client requirements for real-time communication services," "Architectural support for real-time systems: Issues and trade-offs," "Using rate monotonic scheduling technology for real-time communications in a wormhole network," "Priority based real-time communication for large scale wormhole networks," "Simulator for real-time parallel processing architec- tures," "Design and implementation of a priority forwarding router chip for real-time interconnection networks," "Real-time communications scheduling for massively parallel processors," "Providing message delivery guarantees in pipelined flit-buffered multiprocessor networks," "Smart networks for control," "Real-time communication in packet-switched networks," "Real-time communication in multi-hop networks," "Delay jitter control for real-time communication in a packet switching network," "A scheme for real-time channel establishment in wide-area networks," "Rate-controlled service disciplines," "Providing end-to-end performance guarantees using non-work-conserving disciplines," "Effi- cient network QoS provisioning based on per node traffic shap- ing," "The integrated MetaNet architecture: A switch-based multimedia LAN for parallel computing and real-time traffic," "The torus routing chip," "A calculus for network delay, part I: Network elements in isolation," "On the ability of establishing real-time channels in point-to-point packet-switched networks," "Scheduling algorithms for multi-programming in a hard real-time environment," "Virtual new computer communication switching technique," "Virtual-channel flow control," "Hardware support for controlled interaction of guaranteed and best-effort communi- cation," "Support for multiple classes of traffic in multicomputer routers," "PP-MESS-SIM: A flexible and extensible simulator for evaluating multicomputer networks," "Bandwidth requirements for wormhole switches: A simple and efficient design," "The SP2 high-performance switch," "Fast packet switch architectures for broadband integrated services digital networks," "Deadlock-free message routing in multiprocessor interconnection networks," "A survey of wormhole routing techniques in direct networks," "Cost of adaptivity and virtual lanes in a wormhole router," "Routing subject to quality of service constraints in integrated communication networks," "Real-time communication in ATM," "A novel architecture for queue management in the ATM network," "VLSI priority packet queue with inheritance and overwrite," "Exact admission control for networks with bounded delay services," "Hardware-efficient fair queueing architectures for high-speed networks," "Scalable hardware priority queue architectures for high-speed packet switches," "A VLSI sequencer chip for ATM traffic shaper and queue manager," "Systolic priority queues," "Scalable architectures for integrated traffic shaping and link scheduling in high-speed ATM switches," --TR --CTR David Whelihan , Herman Schmit, Memory optimization in single chip network switch fabrics, Proceedings of the 39th conference on Design automation, June 10-14, 2002, New Orleans, Louisiana, USA Sung-Whan Moon , Kang G. Shin , Jennifer Rexford, Scalable Hardware Priority Queue Architectures for High-Speed Packet Switches, IEEE Transactions on Computers, v.49 n.11, p.1215-1227, November 2000 G. Campobello , G. Patan , M. Russo, Hardware for multiconnected networks: a case study, Information SciencesInformatics and Computer Science: An International Journal, v.158 n.1, p.173-188, January 2004 Evgeny Bolotin , Israel Cidon , Ran Ginosar , Avinoam Kolodny, QNoC: QoS architecture and design process for network on chip, Journal of Systems Architecture: the EUROMICRO Journal, v.50 n.2-3, p.105-128, February 2004 Evgeny Bolotin , Israel Cidon , Ran Ginosar , Avinoam Kolodny, Cost considerations in network on chip, Integration, the VLSI Journal, v.38 n.1, p.19-42, October 2004 Kees Goossens , John Dielissen , Jef van Meerbergen , Peter Poplavko , Andrei Rdulescu , Edwin Rijpkema , Erwin Waterlander , Paul Wielage, Guaranteeing the quality of services in networks on chip, Networks on chip, Kluwer Academic Publishers, Hingham, MA,

link scheduling;packet switching;multicomputer router;wormhole switching;real-time communication

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Minimum Achievable Utilization for Fault-Tolerant Processing of Periodic Tasks.

AbstractThe Rate Monotonic Scheduling (RMS) policy is a widely accepted scheduling strategy for real-time systems due to strong theoretical foundations and features attractive to practical uses. For a periodic task set of n tasks with deadlines at the end of task periods, it guarantees a feasible schedule on a single processor as long as the utilization factor of the task set is below n(21/n$-$ 1) which converges to 0.69 for large n. We analyze the schedulability of a set of periodic tasks that is scheduled by the RMS policy and is susceptible to a single fault. The recovery action is the reexecution of all uncompleted tasks. The priority of the RMS policy is maintained even during recovery. Under these conditions, we guarantee that no task will miss a single deadline, even in the presence of a fault, if the utilization factor on the processor does not exceed 0.5. Thus, 0.5 is the minimum achievable utilization that permits recovery from faults before the expiration of the deadlines of the tasks. This bound is better than the trivial bound of 0.69/2 = 0.345 that would be obtained if computation times were doubled to provide for reexecutions in the RMS analysis. Our result provides scheduling guarantees for tolerating a variety of intermittent and transient hardware and software faults that can be handled simply by reexecution. In addition, we demonstrate how permanent faults can be tolerated efficiently by maintaining common spares among a set of processors that are independently executing periodic tasks.

Introduction In the realm of real-time computation, we frequently encounter systems where the tasks are required to execute periodically. Applications where this requirement is common are often found in, for example, process control, space applications, avionics and others. Even when the external events that trigger tasks are not periodic, many real-time systems sample the occurrence of these events periodically and execute the associated tasks during the time slots reserved for them. The sampling rate depends on the expected frequency of the external event. The reason why aperiodic or sporadic tasks are executed in a periodic manner is because the periodic execution is well understood and predictable. A variety of scheduling policies for periodic real-time systems have been studied. A scheduling policy is defined as optimal if it can schedule any feasible set of tasks if any other policy can also do the same. A system is called a fixed-priority system if all the tasks have fixed priorities and these priorities do not change during run time. Rate Monotonic Scheduling (RMS) has been proven to be an optimal scheduling policy for scheduling a set of fixed priority tasks on a uniprocessor. Earliest-Deadline-First (EDF) is the optimal scheduling policy for a variable priority system. Note that a priority of a task is different from its criticality. The former is some measure that is assigned to the tasks by the scheduling policy to facilitate scheduling whereas the latter is the measure of importance of the task as defined by the application. RMS is widely used in practice because it can be easily implemented. It is a preemptive policy where the priority of the tasks are assigned in increasing order of their periods and the task of a particular priority preempts any lower priority task. Liu and Layland proved that as long as the utilization factor of a task set consisting of n tasks is less than n(2 1=n \Gamma 1), the task set is guaranteed a feasible schedule on a uniprocessor [1]. This bound approaches 0:69 as n goes to infinity. However, there may exist task sets which have utilization factors above this bound and still may be feasibly scheduled. The stochastic analysis of the breakdown utilization factor for randomly generated task sets is presented in [2]. The problem for scheduling periodic tasks on multiprocessors is considered in [3] [4] [5]. It is easy to demonstrate that neither the RMS nor the EDF algorithms are optimal for scheduling a set of periodic tasks on a multiprocessor system among fixed and variable priority algorithms respectively [3]. In fact, no scheduling policy is proven to be optimal for a multiprocessor system. Another issue in real-time computing that is currently gaining increased attention of researchers is fault tolerance. Computers are being introduced to a great extent in critical applications and more reliance is being placed on them while reducing human intervention to a minimum. In situations where the demand for hard real-time processing merges with catastrophic consequences of failures, it is not difficult to imagine why fault tolerance must be provided. Responsive systems [6] which must perform computations to successfully meet their deadlines even in the presence of faults are indispensable in many applications. This paper contributes to an evolving framework for the design and implementation of responsive systems. Our goal in this paper is to investigate the issues of fault tolerance in a system of real-time periodic tasks employing Rate Monotonic Scheduling. Previous work has usually addressed software faults where each task has primary and an alternate code. In [7], an off-line scheduling strategy is considered for periodic tasks where the period of a particular task is an integral multiple of the next lower task period. The alternates are scheduled by RMS policy first and then an effort is made to include the maximum number of primary executions in the schedule. A similar problem of scheduling alternate versions of programs called ghosts is considered in [8]. Dynamic programming is used to perform scheduling and an attempt is made to minimize a cost function. A load balancing scheme is presented for periodic task sets scheduled by RMS in [9] where the neighbors of a faulty processor on a ring take over its tasks which are then eventually distributed to the other processors. However, there is no consideration of missing deadlines due to an overload caused by task migration in response to a fault. In this paper, we address the schedulability criterion of a set of periodic tasks for fault-tolerant processing. Specifically, we prove that the minimum achievable utilization is 0.5 for a set of periodic tasks executing in an environment that is susceptible to the occurrence of a single fault where the recovery action is to recompute all the partially executed tasks. This result guarantees that all the tasks will meet their deadlines even in the presence of a fault if the utilization factor of the task set on a processor is less than 0.5. The classes of faults that can be tolerated include intermittent and transient hardware and software faults. In addition, permanent crash and incorrect computation faults can also be handled by providing spares to perform recovery and subsequent execution of the task set. The paper is organized as follows: in Section 2, we provide the background, explain the problem and declare the assumptions. In the following section we present the proof of our assertion that the minimum achievable utilization is 0.5. In Section 4, we address practical and implementation issues. Our conclusions are given in the final section. Background, problem statement and assumptions As has been mentioned in the Introduction, RMS has a strong theoretical foundation and is widely used in practice due to its simplicity. Rate Monotonic Scheduling policy assigns priorities to tasks in the increasing order of their periods. Consider a set S of n tasks. Each Task i is described by a tuple is the execution time of the task, T i is the period and R i is its release time, i.e., the time when the first invocation of the task occurs. Thus ng We will assume that tasks are labeled in such a manner that A task is expected to complete its computation prior to the end of its period. Thus the j th instance (j = 1; 2; .) of the Task i is ready for execution at time R i and has a deadline for completion at We assume that we are dealing with a hard real-time system and the aim is to meet the deadlines under all conditions as opposed to soft real-time systems where the deadlines may be missed and the aim is to reduce the delay. In this paper, when not explicitly mentioned, all R i 's are assumed to be zero. The execution of the tasks is preemptive, i.e., during the execution of a Task i, if any higher priority Task k is ready for execution, the computation of Task i is interrupted and it remains suspended until Task k completes its execution. Then Task i continues from the state at which it was suspended, provided no other task of higher priority is waiting for execution. It is usually assumed that the time to swap tasks is negligible, or that it is accounted for in the computation time. Note that the definition of preemption is recursive, i.e, if Task k has interrupted Task i, it can itself be interrupted by another task of still higher priority. The RMS is a fixed priority policy since the priorities of tasks remain static and do not change during the course of execution of the tasks. The priorities are assigned in the increasing order of the task periods. The task with the smallest period is assigned the highest priority and the task with the largest period the lowest. We will call the arrival time of the task as that instant at which it is ready for execution, i.e., . and its deadline as the next arrival of the same task. The departure time of a task is defined as the time instant when the task completes its execution. Thus the arrival time of the j th instance of Task i is R i its departure time cannot be defined easily because it depends on the parameters of higher priority tasks. The utilization factor U of a task set is defined as For a single processor system, a task set is said to fully utilize the processor under a scheduling algorithm if the task set can be feasibly scheduled using the algorithm but increasing any of the cause the schedule to be infeasible. The least upper bound of the utilization factor is the minimum of the utilization factors for all possible task sets that fully utilize the processor [1] and is also called the minimum achievable utilization [3]. If the task set has a utilization factor which is less than the minimum achievable utilization, then it is guaranteed a feasible schedule. From [1], for a task set with n tasks, the minimum achievable utilization is n(2 1=n \Gamma 1). As n !1, the minimum achievable utilization converges to ln2 which is approximately 0.69. 2.1 Fault classification In any discussion on fault tolerance, it is necessary to consider the issue of fault assumptions because it has a significant impact on the design of the system. Under a crash fault model, the processor is either operating correctly or, if a fault occurs, does not respond at all to any event, internal or external. An incorrect computation fault assumption considers that the processor may fail to produce a correct result in response to correct inputs. For issues related to fault diagnosis and consensus in fault-tolerant processing, the reader can refer to [10]. In addition, faults are also classified as permanent, intermittent and transient [11]. A permanent or hard fault is an erroneous state that is continuous and stable. An intermittent fault occurs occasionally due to unstable nature of hardware. A transient fault results from temporary environmental conditions. A permanent fault can be tolerated only by providing spares which take over the tasks of a primary processor when the fault occurs. Intermittent and transient faults can be tolerated by repeating the computations. 2.2 Analysis of the problem In general, scheduling problem is concerned with allocating shared resources to multiple processes who need the resources simultaneously. This allocation is performed while attempting to achieve certain prespecified goals. In traditional computers, the goal is usually to minimize the total time or increase the response time for all the requests. However, in real-time systems, the goal is simply to allocate the resources in such a manner that the deadlines associated with the tasks are met. In this paper, as we are dealing with scheduling tasks for execution, the resources are the processors. For hard real-time systems, the scheduler has to be such that all tasks are guaranteed to be completed before their deadlines. When real-time systems are to be used for critical applications, it is necessary that the system survives in spite of faults that may arise in the system. Unlike non-real-time systems where the occurrence of faults and subsequent recovery may be permitted to cause delays, it is imperative that the results of computations in real-time systems meet the deadline even in presence of faults. Thus the notion of guaranteeing a feasible schedule has to be extended to cover the random events of fault occurrences. This is a challenging endeavor which has to be addressed nevertheless. In this paper, we will consider fault tolerance strategies for a set of periodic tasks executed under RMS policy which will guarantee that no task will miss even a single deadline due to the occurrence of a fault at any random moment subject to the fault assumptions explicitly stated therein and maintaining the priority of the RMS policy. When one considers introducing fault tolerance into the computation, a host of issues need to be considered in addition to those already existing. The only means of providing fault tolerance is by introducing redundancy in the system. The selection of the appropriate level of time and/or space redundancy is driven by the requirements of the application. Redundancy is provided by creating replicas at some level of computation, usually at the task level in real-time systems. Time redundancy is provided by re-executing the task multiple number of times. The original execution and re-executions can all be performed on a single processor or on different processors. The choice is dependent on the fault model assumption. For real-time systems, time redundancy is the most desirable choice, provided that there is sufficient laxity in the deadlines and there is enough spare capacity that other tasks do not miss their deadlines. This will allow maximum utilization of the available resources. However, if the deadlines are stringent and very little laxity is available, space redundancy is the only choice. Thus an ideal design is one which effectively resolves a tradeoff between these two choices such that minimum cost overhead is incurred and all tasks are guaranteed to meet their deadlines under the fault assumptions. This space-time tradeoff is fundamental to the design of responsive computer systems. The result presented here optimizes the tradeoff to provide scheduling guarantees for a single fault in an environment for periodic tasks. 2.3 Single fault with re-execution of task for recovery We analyze the following scenario: ffl A set of tasks is executing on a single processor and the tasks are scheduled by the RMS policy. ffl All the tasks are independent. ffl A fault may occur at any instant. ffl The interval between successive faults is greater than the largest period in the task set. ffl The fault is detected before the next occurrence of a departure of a task from the processor. For example, if a lower priority tasks is executing during the occurrence of a fault and some time later another higher priority task is supposed to preempt the first task, the fault should be detected before the higher priority task is expected to depart under normal execution. ffl The recovery action is to re-execute all the partially executed tasks at the instant of the fault detection. This includes the currently executing task and all the preempted tasks. ffl The tasks are required to meet their deadlines even if they have to be re-executed due to the occurrence of a fault. ffl The priorities of the RMS policy are maintained even during recovery. Maintaining the priorities of tasks is very important since RMS is a fixed priority scheduling policy and the priorities are assigned at system design time. This approach simplifies the design process because the designer does not have to worry about assigning separate priorities for recovery and analyze the effect of the change in priorities on the schedulability of the task set. One should note that at this stage that we do not place any restrictions on the kind of faults that can be tolerated or the architecture of the system. As long as these conditions are satisfied by the design, the results of this paper are valid. For example, if one were to consider a hardware permanent crash fault, the recovery and subsequent computation would have to be performed on a) Regular execution with no faults. (b) Primary processor, fault occurs just prior to time 17. (c) Spare processor. Figure 1: Feasible schedule in presence of a fault. a spare processor. On the other hand, if a software fault occurs, the recovery is possible on the primary processor itself. An incorrect computation fault can be handled if the fault is detected, perhaps by consistency checks, before the task is expected to depart. In addition, the recovery program for a task need not be the same as the one that is normally executed as long as its computation time is less than or equal to the computation time of the primary code. Two examples are shown in Figures 1 and 2. Both of them consider a task set consisting of two tasks with periods 5 and 7. In these examples, we assume crash faults of processors. In Figure 1, and the processor state as a function of the time is shown under regular execution in Figure 1(a). We observe that the schedule is feasible when no fault occurs. Figures 1(b) and 1(c) show the state of the processor and the spare respectively when a fault occurs just prior to the time instant 17. The fault occurs before Task 2 could complete and so it is re-started on the spare and it meets the deadline of time 21. Figure 2(a) shows the execution profile of the two tasks whose periods are again 5 and 7 respectively. However, in this example 2. Though the schedule is feasible when no fault occurs, the same is not true when a fault causes the recovery action to be taken. The arrival of Task 1 at time 15 preempts Task 2 and a fault occurs just prior to its completion at time 17. So the spare restarts the execution of both tasks, starting with Task 1 as it is a higher priority task. Task 1 completes at time 19 and manages to meet its deadline of time 20. The re-execution of the Task 2 starts at time 19 but is preempted at time 20 by the arrival of the next instance of Task 1 and so Task 2 misses its deadline of time 21. It seems obvious from these examples that certain amount of time redundancy should be provided for recovery and that the RMS scheduling criteria (U ! 0:69) is not sufficient. A trivial a) Regular execution with no faults. (b) Primary processor, fault occurs just prior to time 17. (c) Spare processor. Figure 2: Infeasible schedule in presence of a fault. solution is to "reserve" enough space for all tasks so that in the event of a fault, there is enough spare capacity in terms of time such that the task can be re-executed and still meet its deadline. Since the worst possible time for a fault to occur is just prior to the completion of the task, the amount of extra time to be devoted to task i for recovery is an additional C i . Thus in the Rate Monotonic Analysis of the schedulability of the entire task set, the computation time for all tasks have to be assumed to be 2C i . This means that, in a general case, the effective minimum achievable utilization on each processor is just 0.345, i.e., half of 0.69. However, the situation is not as pessimistic as it appears. We will prove in the following section that a minimum achievable utilization of 0.5 guarantees enough time redundancy to complete recovery before the deadlines. Thus as long as the utilization factor of a task set on a processor is less than or equal to 0.5, the task set is guaranteed a feasible schedule in presence of a single fault. 2.4 Motivation One of the popular traditional approaches to the design of fault-tolerant system is the use of N- modular-redundancy (NMR) [11]. In this technique, every processor is provided with extra spares. The spares may be hot, warm or cold. For real-time systems, hot spares is the preferred choice as no time is wasted to perform recovery. A spare is said to be hot if it synchronously performs all the computations with the primary processor and takes over if the primary processor fails. For fault models such as incorrect computation and Byzantine faults, there may not be any distinction between the primary and the spares as they all perform the same computation and vote on the result to mask faulty results. If we assume crash or fail-stop model, NMR requires that each processor be duplicated to tolerate a single fault and so the number of processors in a fault-tolerant system is 2m where m is the number of processors in the original system. Such a system, called a duplex system, can tolerate up to one fault between the primary and the spare and up to m faults as long as no more than one fault affects a particular primary and its spare. This is achieved by having the space overhead equal to the size of the original system, i.e., by doubling the space resources. The space overhead of duplex system is very high for many applications and it is usually desirable to have a single spare for the group of m processors so that if any processor fails, the spare can be substituted in its place. Whereas providing a single spare is a simple feat in non-real-time systems, ensuring that the recovery will be performed within the deadlines is not easy. The contributions of this paper makes it easy to guarantee recovery by limiting the utilization factor on a processor at 0.5. If U S is the total utilization factor of a large set of tasks, the number of processors needed in a system with a single spare is dU S =0:5e + 1. This assumes crash faults and even distribution of the utilization factor. This is likely to be significantly less than 2 dU S =0:69e in the duplex system. Interestingly, the trivial solution to ensure recovery by doubling the computation time requirements will require dU S =0:35e which is nearly the same as that required by the duplex system. In addition to tolerating hardware crash faults, a major application of the result is towards tolerating software faults. We will deal with this in greater detail in Section 4. 3 Determination of minimum achievable utilization Before we prove that the minimum achievable utilization is 0.5, we present the definitions of some terms used in the proof. The recovery is defined as re-execution of all the partially executed tasks where the priority of the RMS is maintained. Thus during the recovery of a lower priority task, if a higher priority task arrives, the higher priority task will preempt the recovery of the lower priority task. In addition, if the fault affects multiple tasks, higher priority tasks will perform recovery action first. A schedule is said to be feasible for a set of tasks if the task set can be guaranteed a schedule under Rate Monotonic Algorithm (i.e., all tasks will meet their deadlines) even if recovery has to be performed due to a single fault that can occur at any arbitrary instant of time. A set of tasks is said to fully utilize a processor if the task set has a feasible schedule and increasing the computation time of any task in the set causes the schedule to become infeasible. The minimum achievable utilization is the minimum of the utilization factor of every possible sets of tasks that fully utilize the processor. We define a critical instant for a task to be that instant at which an arrival of the task will have the largest response time in the presence of some fault. The schedule of a set of tasks that fully utilizes the processor will have at least one critical instant for some task i where the response time is the period of that task. We shall call that time interval between the arrival and the deadline of the task i as the critical period. A fault that occurs just prior to the completion of a task creates the maximum delay for that task and any lower priority tasks that have been interrupted by it. Hence we only need to examine the effects of a fault at the instants when the tasks are about to be completed. We will consider a number of cases that will lead to the proof of theorem that the minimum achievable utilization is 0.5. 3.1 Case 1: Task set with one task Consider a task set comprising of a single task (C In this case, the release time does not matter. Observation 1 The minimum achievable utilization for a task set with one task is 0.5. Proof: This is obvious since C 1 cannot exceed T 1 =2. If C 1 equals some value x such that occurs at some instant t such that kT 1 44 48 54 55 (a) (b) Recovery of Task 2 Recovery of Task 2 Recovery of Task 1 Figure 3: Schedule of two tasks with periods 6 and 11. then there is not sufficient time to re-execute the task and still meet the deadline at . The processor is fully utilized when C It is important to note that even if the task set has more than one task, the computation time of each of the tasks cannot exceed half the value of its period, i.e., C n is the number of tasks in the set. 3.2 Case 2: Task set with two tasks Before we begin the analysis of the minimum achievable utilization for this case, let us consider the issue of release times. In the traditional RMS analysis the worst delay for the Task 2 is observed when it arrives simultaneously with the Task 1. If the first arrival of the Task 2 can then be feasibly scheduled, any subsequent arrivals will also meet their deadlines and so one has only to consider the feasibility conditions of the simultaneous arrivals of the tasks. This is not necessarily true when one considers the possibility of faults. For example, consider the task set f(1; 6); (4:5; 11)g where release times are zero. By considering just the first arrival, it would appear that the task set has a feasible schedule and the processor is fully utilized. This is shown in Figure 3(a). Tasks 1 and 2 are released simultaneously and since Task 1 has higher priority, it starts execution and departs at begins. A fault occurs just prior to the completion of Task 2 at time instant 5.5 and it is restarted to perform recovery. Task 1 again arrives at time 6 and it preempts recovery. The recovery just completes at time 11 when the next arrival of Task 2 occurs. However, if a fault occurs just before time instant 49, the schedule is infeasible. This is shown in Figure 3(b). Task 2 arrives at time 44 and is preempted by Task 1 which arrives at time 48. A fault occurs just prior to the completion of Task 1 at time 49 and so both tasks have to be re-executed. Task 1 recovers in time at time instant 50 when the recovery of Task 2 begins. However, the next arrival of Task 1 occurs at time 54 and it preempts the recovery of Task 2 and causes it to miss the deadline at time 55. Only 4 units of time are available to Task 2 for recovery in the time interval 50-54 whereas its computation time is 4.5. Thus the correct value of C 2 that fully utilizes the processor is C Hence, in our analysis, we have to consider all possible values of release times. Consider a set of two tasks, arbitrary release times. We will first consider the case when T 2 - 1:5T 1 . Next we will consider various subcases when 3.2.1 Case 2a: Theorem 1 The minimum achievable utilization is 0.5 for a set of two tasks satisfying the condition Proof: We first prove that as long as the utilization factor is less than or equal to 0.5, a feasible schedule is guaranteed for the task set; then we give a particular instance where the processor is fully utilized and the utilization factor is 0.5. From Observation 1, it is clear that C 1 - T 1 =2. Within any interval [R 2 there are at most dT 2 =T 1 e arrivals of Task 1. The worst possible scenario is when Task 2 is about to be completed and is preempted by the arrival of Task 1 and the fault occurs just prior to completion of Task 1. In this case both Tasks 1 and 2 need to be executed again. Task 1 will meet its deadline since C 1 - T 1 =2. Task 2 will meet its deadline if the following condition is i.e. if '- T 2 Under traditional RMS analysis, the feasibility condition is (dT 2 =T 1 But in a fault-tolerant system, each task will have to be executed once more under the worst case scenario of the occurrence of a single fault. Assume that the utilization factor of the task set is less than or equal to 0.5, i.e., Therefore, Thus the feasibility condition given by Equation 2 is guaranteed if Equation 5 is satisfied when T . Thus any task set satisfying the condition of the Theorem 1 is guaranteed a feasible schedule if the utilization factor is less that or equal to 0.5. (a) (b) (c) Figure 4: Modeling subsequent arrivals of the tasks. Now consider the cases when C In each of these two cases, the processor is fully utilized since increasing C 2 in the first case and C 1 in the second case causes the schedule to become infeasible. In both cases, the utilization factor is 0.5. We have also proved that as long as the utilization factor is less than or equal to 0.5, the tasks can be feasibly scheduled. Hence, when T 2 - 1:5T 1 , the minimum achievable utilization is 0.5. 2 3.2.2 Case 2b: We will take the following approach in our proof for this proof: We will first show that each instance of a task can be modeled as the arrival of the first instance with some values of release times R 0and R 0 . Then we will prove that the first instances can be feasibly scheduled for all possible values of release times as long as the utilization factor is less than or equal to 0.5, i.e., we will prove that the minimum achievable utilization among all task sets that fully utilize the processor during the first instance is 0.5. Also, without loss of generality, we can assume that one of R 1 or R 2 is zero Consider Figure 4 where we are interested in the feasibility of meeting the deadline at time instant of the th instance of Task 2 that arrives at time instant R 2 +kT 2 . We consider various cases below where R 1 ffl If R 2 shown in Figure 4(a), the th instance of the Task 2 can be modeled as the first instance of the Task 2 in the task set 0)g. This is possible because any fault during the execution of the (j th instance of Task 1 does not affect the schedulability of the th instance of Task 2. ffl If R 2 shown in Figures 4(b) and (c), the th instance of the Task 2 can be modeled as the first instance of the Task 2 in a task set In the Appendix, we consider all possible cases of the release times and the periods of the tasks. For each of those cases, we present the value of the task computation times that fully utilize the processor during the first instance of the Task 2. For each of those cases, we prove that when the processor is fully utilized during the first instance of Task 2, the utilization factor is greater then 0.5. Theorem 2 The minimum achievable utilization for a set of two tasks satisfying condition 1:5T 1 is 0.5. Proof: We have shown that any subsequent instance of two tasks after the first instance can be modeled as the first instance with some release times. Then we have proved in the Lemmas 3-14 in the Appendix that for all possible values of release times, if the processor is fully utilized for the first instance, the utilization factor is greater than or equal to 0.5. Hence, the minimum achievable utilization for a set of two tasks satisfying condition 3.3 Case 3: Task set with n ? 2 tasks Consider a set of n tasks whose utilization is We will prove by induction that the minimum achievable utilization for a set of n tasks is 0.5. Let us assume that the minimum achievable utilization for a set of tasks is 0.5. We will prove that this is also true for a set of n tasks. Consider the set of the first n \Gamma 1 tasks whose utilization is If both sets S n and S n\Gamma1 have a feasible schedule and U (because Thus we need to consider only those cases where U 0:5. But since U have a feasible schedule because of our assumption. Thus we only need to consider the feasibility of scheduling the Task n. 3.3.1 Case 3a: Theorem 3 The minimum achievable utilization is 0.5 for a set of n tasks satisfying the condition assuming that the minimum achievable utilization of any set of tasks is As in the case of a set of two tasks, if the following condition representing the worst possible scenario is satisfied, the corresponding task set has a feasible schedule. Note that the reverse is not true, i.e., the task set may not satisfy the following condition and still have a feasible schedule. i.e., Assume that U n - 0:5. Therefore Thus the condition in Equation 7 is guaranteed if If Whenever - 0. Thus the sum is also non-negative and Equation 7 is satisfied and the task set is guaranteed a feasible schedule. Thus for all sets of n tasks, the minimum achievable utilization is 0.5 if T n - 1:5T 3.3.2 Case 3b: When we will consider two subcases in the following lemmas. Assume that the set of tasks S n fully utilizes the processor. We note that the set S n\Gamma1 does not fully utilize the processor since U 0:5. Add task n to the set S n\Gamma1 and increment its computation time till the processor is fully utilized and this value of the computation time is C n . Hence only the task n has at least one critical period where the occurrence of a fault and subsequent recovery will cause the task to just its deadline. There are two possible cases: the worst case instant of the occurrence of a fault is just prior to completion of the task n itself in which case the recovery is solely the re-execution of only the task n, or, the worst case instant of occurrence of a fault is just prior to the completion of some other task In the former case, is the fraction of the time that the processor spends executing the task i in the critical period of Task n and x i - dT n =T i e. In the latter case, where y i is the fraction of the time that the processor spends in the normal execution and recovery of the Task i in the critical period of Task n. Here, y Lemma 1 The minimum achievable utilization is 0.5 for a set of n tasks satisfying the conditions e, assuming that the minimum achievable utilization for a set of tasks is 0.5. Construct a set S 0 tasks as follows: The utilization factor U 0 of the set S 0 is the same as that of S n , i.e. U 0 . Now consider a fault just prior to the completion of the task (C during an interval which is a critical interval for the set S n . The time to completion of the task is Thus the last task misses the deadline and so the set S 0 has an infeasible schedule. But since we have assumed that the minimum achievable utilization of a set of tasks 0.5, the utilization factor of S 0 must exceed 0.5. However, U and so U n ? 0:5. Thus the minimum achievable utilization of every set of n tasks that satisfy the conditions of this lemma is 0.5. 2 Here we have proved that every set of n tasks that fully utilizes the processor and satisfies the conditions of the Lemma 1 can be converted into another set of tasks that has an infeasible schedule. As an example, consider a set of three tasks S 0)g. This task set fully utilizes the processor. From this task set, we construct the set S 0 5)g. The set S 0 2 has an infeasible schedule because if a fault occurs at a time just prior to the completion of Task 2 at time instant 2.625, there is not enough spare time to recover. Lemma 2 The minimum achievable utilization is 0.5 for a set of n tasks satisfying the conditions assuming that the minimum achievable utilization of a set of tasks is 0.5. Assume that the set of tasks S n fully utilizes the processor. Since the set S n\Gamma1 does not fully utilize the processor, increment the computation time of the Task n\Gamma1 in S n\Gamma1 so that the utilization factor is 0.5. Let this increase be \Delta and let the new set be S 0 with the utilization factor U 0 \Delta. It is easy to observe that C n - 2\Delta. Since the Task is the lowest priority task in the set of any reduction of the computation time of \Delta from C 0 frees up at least 2\Delta amount of time for Task n that will not be interrupted by the other tasks. The amount is 2\Delta because reduction of \Delta also frees up an extra \Delta from recovery. Thus, ? 0:5We now prove the following theorem for the general case. Theorem 4 For a set of n tasks, the minimum achievable utilization is 0.5. Proof: In Theorem 3 and Lemmas 1 and 2 we have proved that the minimum achievable utilization for a set of n tasks is 0.5 provided that the minimum achievable utilization for a set of tasks is 0.5. In addition, this theorem is true for one task as shown in Observation 1 and has also been proven to be true for a set of two tasks in Theorems 1 and 2. Hence, by induction it is true for all n. 2 Implementation Issues 4.1 Tolerating hardware crash faults Consider a distributed system with a common spare. The spare is not idle but it monitors the state of the processors. After completion of each instance of each task, a processor sends a message to the spare indicating that the task is successfully completed. The spare maintains a list of all tasks in the system and the processor on which they are executing. From this information, it can either be provided a look-up table of all completion times of the tasks or these completion times can be easily computed "on-the-fly". Let Ccomm be the maximum communication latency of the network. If some task was supposed to be completed at time t c , the spare expects a confirmation by the time In case this message is not received, the processor is declared faulty and the spare takes over the faulty processor's task set and initiates recovery. In the rate monotonic analysis of the task set on each processor, the communication time and the overhead in reconfiguration is assumed to be included in the computation time of the task. So, if some task i has computation requirement of C i , then C 0 used for analysis. This technique assumes that the communication delays are finite and bounded, which is not an unreasonable assumption for practical applications. It also requires that the executable code of all tasks be accessible to the spare. As we have discussed in Section 2, the space overhead for guaranteeing deadlines in the presence of a single fault for duplex systems is 2 dU=0:69e processors. However, the number of processors needed for a system with a single spare with recovery is dU=0:5e + 1. U is the total utilization factor of the task set and we assume that the task set is partitioned so that the utilization factor is evenly distributed. Table 1 shows the number of processors required by each scheme for different values of the utilization factors. We observe that providing a common spare significantly reduces the size of the system and the effect is more pronounced for large values of utilization factors. Table 1: Number of processors m in systems where computation times are doubled for RMS analysis, duplex systems and in a system with a common spare for recovery for different values of utilization factor U Doubling computation Duplex system Common spare time in RMS analysis with recovery l U 0:345 l U 0:69 l U 0:5 5 6 19 7 22 22 15 9 28 28 19 100 291 290 200 4.2 Tolerating incorrect computation faults caused by hardware fault Triple Modular Redundancy (TMR) systems are required to tolerate incorrect computation faults. A duplex system can only detect the presence of an incorrect computation fault because the results of the two processors do not agree. A third processor is required so the majority result is assumed to be correct. A similar technique as described above can be used to tolerate a single incorrect computation fault. Rather than having a TMR system, a duplex system with a spare can be used. In case the duplex pair detects an error, the spare is used to perform recovery. The number of processors required for a TMR system is 3 dU=0:69e whereas the number of processors required for a duplex with a spare for recovery is 2 dU=0:5e + 1. Again, U is the total utilization factor for the entire task set. The number of processors required for both schemes is shown in Table 2. We notice that the duplex with a spare again requires less space overhead as compared to a TMR system. However, the benefit is not as large as that observed for crash faults. 4.3 Tolerating software faults and intermittent and transient hardware faults We believe that the greatest application of the results of this paper would be towards tolerating software faults and intermittent and transient hardware faults. In space and hostile industrial applications, outside environment conditions such as alpha particles, electrostatic interference, etc., cause transient errors. In addition software faults such as stack overflows in the operating systems, etc., are best handled by re-execution. By limiting the utilization factor to 0.5 on a processor, we can guarantee that recovery can be performed within the deadlines. Even though we consider the re-execution of all partially executed tasks, it is not necessary if a fault affects a single task. That task can be re-executed to meet its deadline and we can be confident that the re-execution will not cause other tasks to miss their deadlines. In addition, the recovery code need not be the same as the Table 2: Number of processors m in TMR system and in a duplex system with a common spare for recovery for different values of utilization factor U TMR Duplex system with system common spare l U 0:69 l U 0:5 5 9 42 37 100 435 401 primary code. This is especially true for software faults where an alternate program is desirable. As long as the time to execute the recovery program is less than or equal to the execution time of the primary program, we can be certain that the deadlines will be met. 4.4 Tolerating Multiple Faults Multiple faults can be tolerated under our analysis as long as the interval between successive faults is larger than the largest period in the task set. Under this assumption, unlimited transient faults can be tolerated and k permanent crash faults can be tolerated by providing k spares and limiting the utilization factor on each processor to 0.5. For certain task sets and k, NMR system will yield lesser space overhead and greater fault coverage and would be easier to implement. This is the case if 0:5 0:69 where U is the utilization factor for the entire task set. Again we assume that the task set is partitioned so that the utilization factor is evenly distributed. For example, if the total utilization uses only two processors whereas our approach would require three processor. But for most general cases, providing k common spares would result in lesser overheads. Conclusions We have provided a theoretical foundation for fault-tolerant processing of periodic real-time tasks scheduled by the Rate Monotonic Scheduling policy. Under the scenario that recovery from a fault involves restarting all the partially executed tasks while maintaining the priority levels of RMS pol- icy, we show that the minimum achievable utilization on a processor is 0.5. This result guarantees that all tasks will meet their deadlines even in the presence of a fault if the utilization factor on a processor is restricted to 0.5. This bound is much better than the maximum utilization factor of 0.345 (0.69/2) that would be obtained if the computation times of all tasks were naively doubled in RMS analysis to provide for recovery time. The result provides a framework for tolerating transient and intermittent hardware and software faults where re-execution is the preferred recovery technique. In addition, this result is applicable to tolerating permanent crash and incorrect computation faults where spares must be employed to replace faulty processors. In such a system we show that the space redundancy achieved by maintaining a common pool of spares is, in most cases, less than that of an NMR system. The contributions of this paper form an important component in the evolution of Responsive Systems. The concept of providing guarantees of meeting the deadlines in the system in spite of the occurrence of faults is integral to the design of fault-tolerant real-time systems for critical applications. Providing a simple criterion to ensure the feasibility of meeting all deadlines in the presence of a single fault considerably reduces the complexity encountered by designers. This will lead to safer and dependable use of real-time systems for critical applications. --R "Scheduling Algorithms for Multiprogramming in a Hard Real-Time Environment" "The Rate Monotonic Scheduling Algorithm: Exact Characterization and Average Case Behavior" "On a Real-Time Scheduling Problem" "Scheduling Periodically Occurring Tasks on Multiple Proces- sors" "A Note on Preemptive Scheduling of Periodic, Real-Time Tasks" "Responsive Systems: A Challenge for the Nineties" "A Fault-Tolerant Scheduling Problem" "On Scheduling Tasks with a Quick Recovery from Failure" "The Diffusion Model Based Task Remapping for Distributed Real-Time Systems" "The Consensus Problem in Fault-Tolerant Com- puting" The Theory and Practice of Reliable System Design --TR --CTR Rodrigo M. Santos , Jorge Santos , Javier D. Orozco, A least upper bound on the fault tolerance of real-time systems, Journal of Systems and Software, v.78 n.1, p.47-55, October 2005 Sylvain Lauzac , Rami Melhem , Daniel Moss, An Improved Rate-Monotonic Admission Control and Its Applications, IEEE Transactions on Computers, v.52 n.3, p.337-350, March Sasikumar Punnekkat , Alan Burns , Robert Davis, Analysis of Checkpointing for Real-Time Systems, Real-Time Systems, v.20 n.1, p.83-102, Jan. 2001 Frank Liberato , Rami Melhem , Daniel Moss, Tolerance to Multiple Transient Faults for Aperiodic Tasks in Hard Real-Time Systems, IEEE Transactions on Computers, v.49 n.9, p.906-914, September 2000 Tarek F. Abdelzaher , Vivek Sharma , Chenyang Lu, A Utilization Bound for Aperiodic Tasks and Priority Driven Scheduling, IEEE Transactions on Computers, v.53 n.3, p.334-350, March 2004

minimum achievable utilization;real-time systems;rate monotonic scheduling;periodic tasks;fault tolerance

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Efficient Region Tracking With Parametric Models of Geometry and Illumination.

AbstractAs an object moves through the field of view of a camera, the images of the object may change dramatically. This is not simply due to the translation of the object across the image plane. Rather, complications arise due to the fact that the object undergoes changes in pose relative to the viewing camera, changes in illumination relative to light sources, and may even become partially or fully occluded. In this paper, we develop an efficient, general framework for object trackingone which addresses each of these complications. We first develop a computationally efficient method for handling the geometric distortions produced by changes in pose. We then combine geometry and illumination into an algorithm that tracks large image regions using no more computation than would be required to track with no accommodation for illumination changes. Finally, we augment these methods with techniques from robust statistics and treat occluded regions on the object as statistical outliers. Throughout, we present experimental results performed on live video sequences demonstrating the effectiveness and efficiency of our methods.

Introduction Visual tracking has emerged as an important component of systems in several application areas including vision-based control [1, 32, 38, 15], human-computer interfaces [10, 14, 20], surveillance [30, 29, 19], agricultural automation [27, 41], medical imaging [12, 4, 45] and visual reconstruction [11, 42, 48]. The central challenge in visual tracking is to determine the image position of a target region (or features) of an object as it moves through a camera's field of view. This is done by solving what is known as the temporal correspondence problem: the problem of matching the target region in successive frames of a sequence of images taken at closely-spaced time intervals. The correspondence problem for visual tracking has, of course, much in common with the correspondence problems which arise in stereopsis and optical flow. It differs, however, in that the goal is not to determine the exact correspondence for every image location in a pair of images, but rather to determine, in a gross sense, the movement of an entire target region over a long sequence of images. What makes tracking difficult is the extreme variability often present in the images of an object over time. This variability arises from three principle sources: variation in object pose, variation in illumination, and partial or full occlusion of the target. When ignored, any one of these three sources of variability is enough to cause a tracking algorithm to lose its target. In this paper, we develop a general framework for region tracking which includes models for image changes due to motion, illumination, and partial occlusion. In the case of motion, all points in the target region are presumed to be part of the same object allowing us the luxury - at least for most applications - of assuming that these points move coherently in space. This permits us to develop low-order parametric models for the image motion of points within a target region-models that can be used to predict the movement of the points and track the target through an image sequence. In the case of illumination, we exploit the observations of [25, 17, 5] to model image variation due to changing illumination by low-dimensional linear subspaces. The motion and illumination models are then woven together in an efficient algorithm which establishes temporal correspondence of the target region by simultaneously determining motion and illumination parameters. These parameters not only shift and deform image coordinates, but also adjust brightness values within the target region to provide the best match to a fixed reference image. Finally, in the case of partial occlusion, we apply results from robust statistics [16] to show that this matching algorithm is easily extended to include automatic rejection of outlier pixels in a computationally efficient manner. The approach to matching described in this paper is based on comparing the so-called sum-of-squared differences (SSD) between two regions, an idea that has been explored in a variety of contexts including stereo matching [35], optical flow computation [2], hand-eye coordination [38], and visual motion analysis [44]. Much of the previous work using SSD matching for tracking has modeled the motion of the target region as pure translation in the image plane [48, 38], which implicitly assumes that the underlying object is translating parallel to the image plane and is being viewed orthographically. For inter-frame calculations such as those required for optical flow or motion analysis, pure translation is typically adequate. However, for tracking applications in which the correspondence for a finite size image patch must be computed over a long time span, the pure translation assumption is soon violated [44]. In such cases, both geometric image distortions such as rotation, scaling, shear, and illumination changes introduce significant changes in the appearance of the target region and, hence, must be accounted for in order to achieve reliable matching. Attempts have been made to include more elaborate models for image change in region tracking algorithms, but with sizable increases in the computational effort required to establish correspondence. For example, Rehg and Witkin [40] describe energy-based algorithms for tracking deforming image regions, and Rehg and Kanade [39] consider articulated objects undergoing self-occlusion. More recently, Black and Yacoob [8] describe an algorithm for recognizing facial expressions using motion models which include both affine and simple polynomial deformations of the face and its features. Black and Jepson [7] develop a robust algorithm for tracking a target undergoing changes in pose or appearance by combining a simple parametric motion model with an image subspace method [37]. These algorithms require from several seconds to several minutes per frame to compute, and most do not address the problems of changes in appearance due to illumination. In contrast, we develop a mathematical framework for the region tracking problem that naturally incorporates models for geometric distortions and varying illumination. Using this framework, we show that the computations needed to perform temporal matching can be factored to greatly improve algorithm efficiency. The result is a family of region-tracking algorithms which can easily track large image regions (for example the face of a user at a workstation, at a using no special hardware other than a standard digitizer. To date, most tracking algorithms achieving frame-rate performance track only a sparse collection of features (or contours). For example, Blake et al. [9] and Isaard and Blake [33] describe a variety of novel methods for incorporating both spatial and temporal constraints on feature evolution for snake-like contour tracking. Lowe [34] and Gennery [21] describe edge-based tracking methods using rigid three-dimensional geometric models. Earlier work by Ayache [3] and Crowley [13] use incrementally constructed rigid models to constrain image matching. In practice, feature-based and region-based methods can be viewed as complementary techniques. In edge-rich environments such as a manufacturing floor, working with sparse features such as edges has the advantage of computational simplicity - only a small area of the image contributes to the tracking process, and the operations performed in that region are usually very simple. Furthermore, edge-based methods use local derivatives and, hence, tend to be insensitive to global changes in the intensity and/or composition of the incident illumination. However, in less structured situations strong edges are often sparsely distributed in an image, and are difficult to detect and match robustly without a strong predictive model [33]. In such cases, the fact that region-based methods make direct and complete use of all available image intensity information eliminates the need to identify and model a special set of features to track. By incorporating illumination models and robust estimation methods and by making the correspondence algorithm efficient, the robustness and performance of our region tracking algorithms closely rivals that achieved by edge-based methods. The remainder of this article is organized as follows. Section 2 establishes a framework for posing the problem of region tracking for parametric motion models and describes conditions under which an efficient tracking algorithm can be developed. Section 3 then shows how models of illumination can be incorporated with no loss of computational efficiency. Section 4 details modifications for handling partial target occlusion via robust estimation techniques. Section 5 presents experimental results from an implementation of the algorithms. Finally, Section 6 presents a short discussion of performance improving extensions to our tracking algorithm. Tracking Moving Objects In this section, we describe a framework for the efficient tracking of a target region through an image sequence. We first write down a general parametric model for the set of allowable image motions and deformations of the target region. We then pose the tracking problem as the problem of finding the best (in a least squares sense) set of parameter values describing the motions and deformations of the target through the sequence. Finally, we describe how the best set of parameters can be efficiently computed. 2.1 On Recovering Structured Motion We first consider the problem of describing the motion of a target region of an object through a sequence of images. Points on the surface of the object, including those in the target region, are projected down into the image plane. As the object moves through space, the projected points move in the image plane. If the 3-D structure of the object is known in advance, then we could exactly determine the set of possible motions of the points in the images. In general, this information is not known in advance. Therefore, we approximate the set of possible motions by a parametric model for image motions. Let I(x; t) denote the brightness value at the location in an image acquired at time t and let r x I(x; t) denote the spatial gradient at that location and time. The symbol denotes an identified "initial" time and we refer to the image at time t 0 as the reference image. Let the set be a set of N image locations which define a target region. We refer to the brightness values of the target region in the reference image as the reference template. Over time, the relative motion between the target object and the camera causes the image of the target to shift and to deform. Let us model the image motion of the target region of the object by a parametric motion model f(x; -) parameterized by - with We assume that f is differentiable in both - and x: We call - the motion parameter vector. We consider recovering the motion parameter vector for each image in the tracking sequence as the equivalent to "tracking the object." We write - (t) to denote the ground truth values of these parameters at time t, and -(t) to denote the corresponding estimate. The argument t will be suppressed when it is obvious from its context. Suppose that a reference template is acquired at time t 0 and that initially - us assume for now that the only changes in subsequent images of the target are completely described by f ; i.e. there are no changes in the illumination of the target. It follows that for any time t ? t 0 there is a parameter vector - (t) such that This a generalization of the so-called image constancy assumption [28]. The motion parameter vector of the target region can be estimated at time t by minimizing the following least squares objective function For later developments, it is convenient to rewrite this optimization problem in vector notation. To this end, let us consider images of the target region as vectors in an N dimensional space. The image of the target region at time t, under the change of coordinates f with parameters -; is written as This vector is subsequently referred to as the rectified image at time t with parameters -: We also make use of the partial derivatives of I with respect to the components of - and the time parameter t: These are written as I - i I - i I - i I - i and I t (-; @t I I t (f(x I t (f(x N ; -); t) Using this vector notation, the image constancy assumption (1) can be rewritten as I(- (t); and (2) becomes In general, (6) is a non-convex objective function. Thus, in the absence of a good starting point, this problem will usually require some type of costly global optimization procedure to solve [6]. In the case of visual tracking, the continuity of motion provides such a starting point. Suppose that, at some arbitrary time the geometry of the target region is described by -(t): We recast the tracking problem as one of determining a vector of offsets, ffi- such that -(t ffi- from an image acquired at t + -: Incorporating this modification into (6), we redefine the objective function as a function on ffi- If the magnitude of the components of ffi- are small, then it is possible to apply continuous optimization procedures to a linearized version of the problem [7, 28, 35, 47, 44]. The linearization is carried out by expanding I(- in a Taylor series about - and t; where h:o:t denotes higher order terms of the expansion, and M is the Jacobian matrix of I with respect to -; i.e. the N \Theta n matrix of partial derivatives which can be written in column form as While the notation above explicitly indicates that the values of the partial derivatives are a function of the evaluation point (-; t); these arguments will be suppressed when obvious from their context. By substituting (8) into (7) and ignoring the higher order terms, we have With the additional approximation -I t (-; O(ffi- Solving the set of equations yields the solution provided the matrix M is full rank. When this is not the case, we are faced with a generalization of the aperture problem, i.e. the target region does not have sufficient structure to determine all of the elements of - uniquely. In subsequent developments, it will be convenient to define the error vector Incorporating this definition into (12), we see that the solution of (6) at time t a solution at time t is where M is evaluated at (-; t): 2.2 An Efficient Tracking Algorithm From (13), we see that to track the target region through the image sequence, we must compute the Jacobian matrix M(-; t): Each element of this matrix is given by where r f I is the gradient of I with respect to the components of the vector f . Recall that the Jacobian matrix of the transformation f regarded as a function of - is the 2 \Theta n matrix By making use of (15), M can be written compactly in row form as Because M depends on time-varying quantities, it may appear that it must be completely recomputed at each time step-a computationally expensive procedure involving the calculation of the image gradient vector, the calculation of a 2 \Theta n Jacobian matrix, and n 2 \Theta 1 vector inner products for each of the N pixels of the target region. However, we now show that it is possible to reduce this computation by both eliminating the need to recompute image gradients and by factoring M: First, we eliminate the need to compute image gradients. To do so, let us assume that our estimate is exact, i.e. differentiating both sides of (1) we obtain is the 2 \Theta 2 Jacobian matrix of f treated as a function of @x @y Combining (17) with (16), we see that M can be written as r x I(x It follows that for any choice of image deformations, the image spatial gradients need only be calculated once on the reference template. This is not surprising given that the target at only a distortion of the target at time t 0 ; and so its image gradients are also a distortion of those at t 0 : This transformation also allows us to drop the time argument of M and regard it solely as a function of -: The remaining non-constant factor in M is a consequence of the fact that, in general, f x and f - involve components of - and, hence, implicitly vary with time. However, suppose that we choose f so that f \Gamma1 can be factored into the product of a 2 \Theta k matrix \Gamma which depends only on image coordinates, and a k \Theta n matrix \Sigma which depends only on - as For example, as discussed in more detail below, one family of such factorizations results when f is a linear function of the image coordinate vector x: Combining (19) with (20), we have r x r x I(x As a result, we have shown that M can be written as a product of an constant N \Theta k matrix M 0 and a time-varying k \Theta n matrix \Sigma: We can now exploit this factoring to define an efficient tracking algorithm which operates as follows: offline: ffl Define the target region. ffl Acquire and store the reference template. ffl Compute and store M 0 and online: ffl Use the most recent motion parameter estimate -(t) to rectify the target region in the current image. by taking the difference between the rectified image and the reference template. ffl Solve the system \Sigma T \Sigma is evaluated at The online computation performed by this algorithm is quite small, and consists of two n \Theta k matrix multiplies, k N-vector inner products, n k-vector inner products, and an n \Theta n linear system solution, where k and n are typically far smaller than N: We note that the computation can be further reduced if \Sigma is invertible. In this case, the solution to the linear system can be expressed as where \Sigma is evaluated at -(t): The factor (M 0 T can be computed offline, so the online computation is reduced to n N-vector inner products and n n-vector inner products. 2.3 Some Examples 2.3.1 Linear Models Let us assume that f(x; -) is linear in x: Then we have and, hence, f x It follows that f \Gamma1 x f - is linear in the components of x and the factoring defined in (20) applies. We now present three examples illustrating these concepts. Pure Translation: In the case of pure translation, the allowed image motions are parameterized by the vector It follows immediately that f x and f - are both the 2 \Theta 2 identity matrix, and therefore and \Sigma is the 2 \Theta 2 identity matrix. The resulting linear system is nonsingular if the image gradients in the template region are not all collinear, in which case the solution at each time step is just Note that in this case constant matrix which can be computed offline. Translation, Rotation and Scale: Objects which are viewed under scaled orthography and which do not undergo out-of-plane rotation can be modeled using a four parameter model consisting an image-plane rotation through an angle '; a scaling by s; and a translation by u: The change of coordinates is given by where R(') is a 2 \Theta 2 rotation matrix. After some minor algebraic manipulations, we obtain and From this M 0 can be computed using (21) and, since \Sigma is invertible, the solution to the linear system becomes This result can be explained as follows. The matrix M 0 is the linearization of the system about the target has orientation '(t) and s(t): Image rectification effectively rotates the target by \Gamma' and scales by 1 s so the displacements of the target are computed in the original target coordinate system. \Sigma \GammaT then applies a change of coordinates to rotate and scale the computed displacements from the original target coordinate system back to the actual target coordinates. Affine Motion: The image distortions of planar objects viewed under orthographic projection are described by a six-parameter linear change of coordinates. Suppose that we define a c b d After some minor algebraic manipulations, we obtain and Note that \Sigma is once again invertible which allows for additional computational savings as before. 2.3.2 Nonlinear Motion Models The separability property needed for factoring does not hold for any type of nonlinear motion. However, consider a motion model of the form Intuitively, this model performs a quadratic distortion of the image according to the equation example, a polynomial model of this form was used in [8] to model the motions of lips and eyebrows on a face. Again, after several algebraic steps we arrive at and Note this general result holds for any distortion which can be expressed exclusively as either adding more freedom to the motion model, for example combining affine and polynomial distortion, often makes factoring impossible. One possibility in such cases is to use a cascaded model in which the image is first rectified using an affine distortion model, and then the resulting rectified image is further rectified for polynomial distortion. 2.4 On the Structure of Image Change The Jacobian matrix M plays a central role in the algorithms described above, so it is informative to digress briefly on its structure. If we consider the rectified image as a continuous time-varying quantity, then its total derivative with respect to time is dI dt dt I t or - Note that this is simply a differential form of (8). Due to the image constancy assumption (1), it follows that - This is, of course, a parameterized version of Horn's optical flow constraint equation [28]. In this form, it is clear that the role of M is to relate variations in motion parameters to variations in brightness values in the target region. The solution given in (13) effectively reverses this relationship and provides a method for interpreting observed changes in brightness as motion. In this sense, we can think of the algorithm as performing correlation on temporal changes (as opposed to spatial structure) to compute motion. To better understand the structure of M; recall that in column form, it can be written in terms of the partial derivatives of the rectified image: Thus, the model states that the temporal variation in image brightness in the target region is a weighted combination of the vectors I - i : We can think of each of these columns (which have an entry for every pixel in the target region) as a "motion template" which directly represents the changes in brightness induced by the motion represented by the corresponding motion parameter. For example, in the top row of Figure 1, we have shown these templates for several canonical motions of an image of a black square on a white background. Below, we show the corresponding templates for a human face. The development in this section has assumed that we start with a given parametric motion model from which these templates are derived. Based on that model, the structure of each entry of M is given by (15) which states that The image gradient r f I defines, at each point in the image, the direction of strongest intensity change. The vector f - j evaluated at x i is the instantaneous direction and magnitude of motion of that image location captured by the parameter The collection of the latter for all pixels in the region represents the motion field defined by the motion parameter Target Image X Translation Y Translation Rotation Scale Target Image X Translation Y Translation Rotation Scale Figure 1: Above, the reference template for a bright square on a dark background the motion template for four canonical motions. Below, the same motion templates for a human face. Thus, the change in the brightness of the image location x i due to the motion parameter - j is the projection of the image gradient onto the motion vector. This suggests how our techniques can be used to perform structured motion estimation without an explicit parametric motion model. First, if the changes in images due to motion can be observed directly (for example, by computing the differences of images taken before and after small reference motions are performed), then these can be used as the motion templates which comprise M: Second, if a one or more motion fields can be observed (for example, by tracking a set of fiducial points in a series of training images), then projecting each element of the motion field onto the corresponding image gradient yields motion templates for those motion fields. The linear estimation process described above can be used to time-varying images in terms of those basis motions. Illumination-Insensitive Tracking The systems described above are inherently sensitive to changes in illumination of the target region. This is not surprising, as the incremental estimation step is effectively computing a structured optical flow, and optical flow methods are well-known to be sensitive to illumination changes [28]. Thus, shadowing or shading changes of the target object over time lead to bias, or, in the worst case, complete loss of the target. Recently, it has been shown that a relatively small number of "basis" images can often be used to account for large changes in illumination [5, 17, 22, 24, 43]. Briefly, the reason for this is as follows. Consider a point p on a Lambertian surface and a collimated light source characterized by a vector s 2 IR 3 , such that the direction of s gives the direction of the light rays and ksk gives the intensity of the light source. The irradiance at the point p is given by where n is the unit inwards normal vector to the surface at p and a is the non-negative absorption coefficient (albedo) of the surface at the point p [28]. This shows that the irradiance at the point p, and hence the gray level seen by a camera, is linear on s 2 IR 3 . Therefore, in the absence of self-shadowing, given three images of a Lambertian surface from the same viewpoint taken under three known, linearly independent light source di- rections, the albedo and surface normal can be recovered; this is the well-known method of photometric stereo [50, 46]. Alternatively, one can reconstruct the image of the surface under a novel lighting direction by a linear combination of the three original images [43]. In other words, if the surface is purely Lambertian and there is no shadowing, then all images under varying illumination lie within a 3-D linear subspace of IR N , the space of all possible images (where N is the number of pixels in the images). A complication comes when handling shadowing: all images are no longer guaranteed to lie in a linear subspace [5]. Nevertheless, as done in [24], we can still use a linear model as an approximation: a small set of basis images can account for much of the shading changes that occur on patches of non-specular surfaces. Naturally, we need more than three images (we use between 8 and 15) and a higher than three dimensional linear subspace (we use 4 or if we hope to provide good approximation to these effects. Returning to the problem of region tracking, suppose now that we have a basis of image vectors where the ith element of each of the basis vectors corresponds to the image location x us choose the first basis vector to be the template image, To model brightness changes, let us choose the second basis vector to be a column of ones, i.e. us choose the remaining basis vectors by In practice, choosing a value close to the mean of the brightness of the image produces a more numerically performing SVD (singular value decomposition) on a set of training images of the target, taken under varying illumination. We denote the collection of basis vectors by the matrix Suppose now that so that the template image and the current target region are registered geometrically at time t: The remaining difference between them is due to illumination. From the above discussion, it follows that interframe changes in the current target region can be approximated by the template image plus a linear combination of the basis vectors B, i.e. where the vector - . Note that because the template image and an image of ones are included in the basis B, we implicitly handle both variation due to contrast changes and variation due to brightness changes. The remaining basis vectors are used to handle more subtle variation - variation that depends both on the geometry of the target object and on the nature of the light sources. Using the vector-space formulation for motion recovery established in the previous sec- tion, it is clear that illumination and geometry can be recovered in one global optimization step solved via linear methods. Incorporating illumination into (7) we have the following modified optimization: Substituting (42) into (43) and performing the same simplifications and approximations as before, we arrive at Solving ffi- In most tracking applications, we are only interested in the motion parameters. We can eliminate explicit computation of these parameters by first optimizing over - in (44). Upon stable linear system. substituting the resulting solution back into (44) and then solving for ffi- we arrive at Note that if the columns of B are orthogonal vectors, B T B is the identity matrix. It is easy to show that in both equations, factoring M into time-invariant and time-varying components as described above leads to significant computational savings. Since the illumination basis is time-invariant, the dimensionality of the time-varying portion of the computation depends only on the number of motion fields to be computed, not on the illumination model. Hence, we have shown how to compute image motion while accounting for variations in illumination using no more online computation than would be required to compute pure motion. Making Tracking Resistant to Occlusion As a system tracks objects over a large space, it is not uncommon that other objects "intrude" into the picture. For example, the system may be in the process of tracking a target region which is the side of a building when, due to observer motion, a parked car begins to occlude a portion of that region. Similarly the target object may rotate, causing the tracked region to "slide off" and pick up a portion of the background. Such intrusions will bias the motion parameter estimates and, in the long term can potentially cause mistracking. In this section, we describe how to avoid such problems. For the sake of simplicity, we develop a solution for the case where we are only recovering motion parameters; the modifications for combined motion and illumination models are straightforward. A common approach to this problem is to assume that occlusions create large image differences which can be viewed as "outliers" by the estimation process [7]. The error metric is then modified to reduce sensitivity to "outliers" by solving a robust optimization problem of the form where ae is one of a variety of "robust" regression metrics [31]. It is well-known that optimization of (47) is closely related to another approach to robust estimation-iteratively reweighted least squares (IRLS). We have chosen to implement the optimization using a somewhat unusual form of IRLS due to Dutter and Huber [16]. In order to formulate the algorithm, we introduce the notation of an "inner iteration" which is performed one or more times at each time step. We will use a superscript to denote this iteration. Let ffi- i denote the value of ffi- computed by the ith inner iteration with ffi- the vector of residuals in the ith iteration r i as We introduce a diagonal weighting matrix W which has entries The inner iteration cycle at time t + - is consists of performing an estimation step by solving the linear system where \Sigma is evaluated at -(t)) and r i and W i are given by (48) and (49), respectively. This process is repeated for k iterations. This form of IRLS is particularly efficient for our problem. It does not require recomputation of or \Sigma and, since the weighting matrix is diagonal, does not add significantly to the overall computation time needed to solve the linear system. In addition, the error vector e is fixed over all inner iterations, so these iterations do not require the additional overhead of acquiring and warping images. As discussed in [16], on linear problems this procedure is guaranteed to converge to a unique global minimum for a large variety of choices of ae: In this article, ae is taken to be a so-called "windsorizing" function [31] which is of the form where r is normalized to have unit variance. The parameter - is a user-defined threshold which places a limit on the variations of the residuals before they are considered outliers. This function has the advantage of guaranteeing global convergence of the IRLS method while being cheap to compute. The updating function for matrix entries is As stated, the weighting matrix is computed anew at each iteration, a process which can require several inner iterations. However, given that tracking is a continuous process, it is natural to start with an initial weighting matrix that is closely related to that computed at the end of the previous estimation step. In doing so, two issues arise. First, the fact that the linear system we are solving is a local linearization of a nonlinear system means that, in cases when inter-frame motion is large, the effect of higher-order terms of the Taylor series expansion will cause areas of the image to masquerade as outliers. Second, if we assume that areas of the image with low weights correspond to intruders, it makes sense to add a "buffer zone" around those areas for the next iteration to pro-actively cancel the effects of intruder motion. Both of these problems can be dealt with by noting that the diagonal elements of W themselves form an image where "dark areas" (those locations with low value ) are areas of occlusion or intrusion, while "bright areas" (those with value 1) are the expected target. Let Q(x) to be the pixel values in the eight-neighborhood of the image coordinate x plus the value at x itself. We use two common morphological operators [26] and When applied to a weighting matrix image, close has the effect of removing small areas of outlier pixels, while open increases their size. Between frames of the sequence we propagate the weighting matrix forward after applying one step of close to remove small areas of outliers followed by two or three steps of open to buffer detected intruders. 5 Implementation and Experiments This section illustrates the performance of the tracking algorithm under a variety of circum- stances, noting particularly the effects of image warping, illumination compensation, and outlier detection. All experiments were performed on live video sequences by an SGI Indy equipped with a 175Mhz R4400 SC processor and VINO image acquisition system. Rotation Scale Aspect Ratio Shear Figure 2: The columns of the motion Jacobian matrix for the planar target and their geometric interpretations. 5.1 Implementation We have implemented the methods described above within the X Vision environment [23]. The implemented system incorporates all of the linear motion models described in Section 2, non-orthonormal illumination bases as described in Section 3, and outlier rejection using the algorithm described in Section 4. The image warping required to support the algorithm is implemented by factoring linear transformations into a rotation matrix and a positive-definite upper-diagonal matrix. This factoring allows image warping to be implemented in two stages. In the first stage, an image region surrounding the target is acquired and rotated using a variant on standard Bresenham line-drawing algorithms [18]. The acquired image is then scaled and sheared using a bilinear interpolation. The resolution of the region is then reduced by averaging neighboring pixels. Spatial and temporal derivatives are computed by applying Prewitt operators on the reduced scale images. More details on this level of the implementation can be found in [23]. Timings of the algorithm 2 indicate that it can perform frame rate (30 Hz) tracking of image regions of up to 100 \Theta 100 pixels at one-half resolution undergoing affine distortions and illumination changes. Similar performance has been achieved on a 120Mhz Pentium processor and 70 Mhz Sun SparcStation. Higher performance is achieved for smaller regions, lower resolutions, or fewer parameters. For example, tracking the same size region while computing just translation at one-fourth resolution takes just 4 milliseconds per cycle. 5.2 Planar Tracking As a baseline, we first consider tracking a non-specular planar object-the cover of a book. Affine warping augmented with brightness and contrast compensation is the best possible linear approximation to this case (it is exact for an orthographic camera model and purely Lambertian surface). As a point of comparison, recent work by Black and Jepson [7] used the rigid motion plus scaling model for SSD-based region tracking. Their reduced model is more efficient and may be more stable since fewer parameters must be computed, but it does ignore the effects of changing aspect ratio and shear. We tested both the rigid motion plus scale (RM+S) and full affine motion models on the same live video sequence of the book cover in motion. Figure 2 shows the set of motion templates (the columns of the motion matrix) for an 81 \Theta 72 region of a book cover tracked at one third resolution. Figure 3 shows the results of tracking. The upper series of images shows several images of the object with the region tracked indicated with a black frame (the RM+S algorithm) and a white frame (the FA algorithm). The middle row of images shows the output of the warping operator from the RM+S algorithm. If the computed parameters were error-free, these images would be identical. However, because of the inability to correct for aspect ratio and skew, the best fit leads to a skewed image. The bottom row shows the output of the warping operator for the FA algorithm. Here we see that the full affine warping is much better at accommodating the full range of image distortions. The graph at the bottom of the figure shows the least squares residual (in squared gray-values per pixel). Here, the difference between the two geometric models is clearly evident. 5.3 Human Face Tracking There has been a great deal of recent interest in face tracking in the computer vision literature [8, 14, 36]. Although faces can produce images with significant variation due to empirical results suggest that a small number of basis images of a face gathered under different illuminations is sufficient to accurately account for most gross shading and illumination effects [24]. At the same time, the depth variations exhibited by facial features are small enough to be well-approximated by an affine warping model. The following 2 Because of additional data collection overhead, the tracking performance in the experiments presented here is slower than the stated figures. Frame 0 Frame 50 Frame 70 Frame 120 Frame 150 Frame 230 Residuals: Planar Test FA RM+S Gray values Frames10.0020.0030.0040.00 Figure 3: Top, several images of a planar region and the corresponding warped image computed by a tracker computing position, orientation and scale (RM+S), and one computing a full affine deformation (FA). The image at the left is the initial reference image. Bottom, the graph of the SSD residuals for both algorithms. experiments demonstrate the ability of our algorithm to track a face as it undergoes changes in pose and illumination, and under partial occlusion. Throughout, we assume the subject is roughly looking toward the camera, so we use the rigid motion plus scaling (RM+S) motion model. Figure 1 on page 14 shows the columns of the motion matrix for this model. 5.3.1 Geometry We first performed a test to determine the accuracy of the computed motion parameters for the face and to investigate the effect of the illumination basis on the sensitivity of those estimates. During this test, we simultaneously executed two tracking algorithms: one using the rigid motion plus scale model (RM+S) and one which additionally included an illumination model for the face (RM+S+I). The algorithms were executed on a sequence which did not contain large changes in the illumination of the target. The top row of Figure 4 shows images excerpted from the video sequence. In each image, the black frames denote the region selected as the best match by RM+S and the white frames correspond to the best match computed by RM+S+I. For this test, we would expect both algorithms to be quite accurate and to exhibit similar performance unless the illumination basis significantly affected the sensitivity of the computation. As is apparent from the figures, the computed motion parameters of both algorithms are extremely similar for the entire run - so close that in many cases one frame is obscured by the other. In order to demonstrate the absolute accuracy of the tracking solution, below each live image in Figure 4 we have included the corresponding rectified image computed by RM+S+I. The rectified image at time 0 is the reference template. If the motion of the target fit the RM+S motion model, and the computed parameters were exact, then we would expect each subsequent rectified image to be identical to the reference template. Despite the fact that the face is non-planar and we are using a reduced motion model, we see that the algorithm is quite effective at computing an accurate geometric match. Finally, the graph in Figure 4 shows the residuals of the linearized SSD computation at each time step. As is apparent from the figures, the residuals of both algorithms are also extremely similar for the entire run. From this experiment we conclude that, in the absence of illumination changes, the performance of both algorithms is quite similar - including illumination models does not appear to reduce accuracy. Frame 0 Frame Residuals: Face with No Lighting Changes RM+S+I RM+S Gray values Frames20.00 Figure 4: Top row, excerpts from a sequence of tracked images of a face. The black frames represent the region tracked by an SSD algorithm using no illumination model (RM+S) and the white frames represent the regions tracked by an algorithm which includes an illumination model (RM+S+I). In some cases the estimates are so close that only one box is visible. Middle row, the region within the frame warped by the current motion estimate. Bottom row, the residuals of the algorithms expressed in gray-scale units per pixel as a function of time. Figure 5: The illumination basis for the face (B). The left two images are included to compensate for brightness and contrast, respectively, while the remaining four images compensate for changes in lighting direction. 5.3.2 Illumination In a second set of experiments, we kept the face nearly motionless and varied the illumination. We used an illumination basis of four orthogonal image vectors. This basis was computed offline by acquiring ten images of the face under various lighting conditions. A singular value decomposition (SVD) was applied to the resulting image vectors and the vectors with the maximum singular values were chosen to be included in the basis. The illumination basis is shown in Figure 5. Figure 6 shows the effects of illumination compensation for the illumination situations depicted in the first row. As with warping, if the compensation were perfect, the images of the bottom row would appear to be identical up to brightness and contrast. In particular, note how the strong shading effects of frames 70 through 150 have been "corrected" by the illumination basis. 5.3.3 Combining Illumination and Geometry Next, we present a set of experiments illustrating the interaction of geometry and illumi- nation. In these experiments we again executed two algorithms again labeled RM+S and RM+S+I. As the algorithms were operating, a light was periodically switched on and off and the face moved slightly. The results appear in Figure 7. In the residual graph, we see that the illumination basis clearly "accounts" for the shading on the face quite well, leading to a much lower fluctuation of the residuals. The sequence of images shows an excerpt near the middle of the sequence where the RM+S algorithm (which could not compensate for il- 0Figure The first row of images shows excerpts of a tracking sequence. The second row is a magnified view of the region in the white frame. The third row contains the images in the second row after adjustment for illumination using the illumination basis shown in Figure 5 (for sake of comparison we have not adjusted for brightness and contrast). Frame 90 Frame 100 Frame 110 Frame 120 Frame 130 Frame 140 Frame 150 Residuals: Face with Illumination Changes RM+S+I RM+S Gray values Frames20.000.00 50.00 100.00 150.00 200.00 250.00 300.00 Figure 7: Top, an excerpt from a tracking sequence containing changes in both geometry and illumination. The black frame corresponds to the algorithm without illumination (RM+S) and the write frame corresponds to the algorithm with an illumination basis (RM+S+I). Note that the algorithm which does not use illumination completely looses the target until the original lighting is restored. Bottom, the residuals, in gray scale units per pixel, of the two algorithms as a light is turned on and off. lumination changes) completely lost the target for several frames, only regaining it after the original lighting was restored. Since the target was effectively motionless during this period, this can be completely attributed to biases due to illumination effects. Similar sequences with larger target motions often cause the purely geometric algorithm to loose the target completely. 5.3.4 Tracking With Outliers Finally, we illustrate the performance of the method when the image of the target becomes partially occluded. We again track a face. The motion and illumination basis are the same as before. In the weighting matrix calculations the pixel variance was set to 5 and the outlier threshold was set to 5 variance units. The sequence is an "office" sequence which includes several "intrusions" including the background, a piece of paper, a telephone, a soda can, and a hand. As before we executed two versions of the tracker, the non-robust algorithm from the previous experiment (RM+S+I) and a robust version (RM+S+I+O). Figure 8 shows the results. The upper series of images shows the region acquired by both algorithms (the black frame corresponds to RM+S+I, the white to RM+S+I+O). As is clear from the sequence, the non-robust algorithm is disturbed significantly by the occlusion, whereas the robust algorithm is much more stable. In fact, a slight motion of the head while the soda can is in the image caused the non-robust algorithm to mistrack completely. The middle series of images shows the output of the warping operation for the robust algorithm. The lower row of images depicts the weighting values attached to each pixel in the warped image. Dark areas correspond to "outliers." Note that, although the occluded region is clearly identified by the algorithm, there are some small regions away from the occlusion which received a slightly reduced weight. This is due to the fact that the robust metric used introduces some small bias into the computed parameters. In areas where the spatial gradient is large (e.g. near the eyes and mouth), this introduces some false rejection of pixels. It is also important to note that the dynamical performance of the tracker is significantly reduced by including outliers. Large, fast motions tend to cause the algorithm to "turn off" areas of the image where there are large gradients, slowing convergence. At the same time, performing outlier rejection is more computationally intensive as it requires explicit computation of both the motion and illumination parameter to calculate the residual values. 6 Discussion and Conclusions We have shown a straightforward and efficient solution to the problem of tracking regions undergoing geometric distortion, changing illumination, and partial occlusion. The method is simple, yet robust, and it builds on an already popular method for solving spatial and temporal correspondence problems. Although the focus in this article has been on parameter estimation techniques for tracking using image rectification, the same estimation methods can be used for directly controlling devices. For example, instead of computing a parameter estimate -; the incremental solutions ffi- can be used to control the position and orientation of a camera so to stabilize the target image by active motion. Hybrid combinations of camera control and image warping are also possible. Frame 0 Frame Residuals: Face with Partial Occlusion RM+S+I+O RM+S+I gray values Figure 8: The first row of images shows excerpts of a tracking sequence with occurrences of partial occlusion. The black frame corresponds to the algorithm without outlier rejection (RM+S+I) and the write frame corresponds to the algorithm with outlier rejection (RM+S+I+O). The second row is a magnified view of the region in the white frame. The third row contains the corresponding outlier images where darker areas mark outliers. The graph at the bottom compares the residual values for both algorithms. One possible objection to the methods is the requirement that the change from frame to frame is small, limiting the speed at which objects can move. Luckily, there are several means for improving the dynamical performance of the algorithms. One possibility is to include a model for the motion of the underlying object and to incorporate prediction into the tracking algorithm. Likewise, if a model of the noise characteristics of images is available, the updating method can modified to incorporate this model. In fact, the linear form of the solution makes it straightforward to incorporate the estimation algorithm into a Kalman filter or similar iterative estimation procedure. Performance can also be improved by operating the tracking algorithm at multiple levels of resolution. One possibility, as is used by many authors [7, 44], is to perform a complete coarse to fine progression of estimation steps on each image in the sequence. Another possi- bility, which we have used successfully in prior work [23], is to dynamically adapt resolution based on the motion of the target. That is, when the target moves quickly estimation is performed at a coarse resolution, and when it moves slowly the algorithm changes to a higher resolution. The advantage of this approach is that it not only increases the range over which the linearized problem is valid, but it also reduces the computation time required on each image when motion is fast. We are actively continuing to evaluate the performance of these methods, and to extend their theoretical underpinnings. One area that still needs attention is the problem of determining an illumination basis online, i.e. while tracking the object. Initial experiments in this direction have shown that online determination of the illumination basis can be achieved, although we have not included such results in this paper. As in [7], we are also exploring the use of basis images to handle changes of view or aspect not well addressed by warping. We are also looking at the problem of extending the method to utilize shape information on the target when such information is available. In particular, it is well known [49] that under orthographic projection, the image deformations of a surface due to motion can be described with a linear motion model. This suggests that our methods can be extended to handle such models. Furthermore, as with the illumination basis, it may be possible to estimate the deformation models online, thereby making it possible to efficiently track arbitrary objects under changes in illumination, pose, and partial occlusion. Acknowledgments This research was supported by ARPA grant N00014-93-1-1235, Army DURIP grant DAAH04-95- 1-0058, National Science Foundation grant IRI-9420982, Army Research Office grant DAAH04-95- 1-0494, and by funds provided by Yale University. 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motion estimation;real-time vision;visual tracking;illumination;robust statistics

290124

Scale-Space Derived From B-Splines.

AbstractIt is well-known that the linear scale-space theory in computer vision is mainly based on the Gaussian kernel. The purpose of the paper is to propose a scale-space theory based on B-spline kernels. Our aim is twofold. On one hand, we present a general framework and show how B-splines provide a flexible tool to design various scale-space representations: continuous scale-space, dyadic scale-space frame, and compact scale-space representation. In particular, we focus on the design of continuous scale-space and dyadic scale-space frame representation. A general algorithm is presented for fast implementation of continuous scale-space at rational scales. In the dyadic case, efficient frame algorithms are derived using B-spline techniques to analyze the geometry of an image. Moreover, the image can be synthesized from its multiscale local partial derivatives. Also, the relationship between several scale-space approaches is explored. In particular, the evolution of wavelet theory from traditional scale-space filtering can be well understood in terms of B-splines. On the other hand, the behavior of edge models, the properties of completeness, causality, and other properties in such a scale-space representation are examined in the framework of B-splines. It is shown that, besides the good properties inherited from the Gaussian kernel, the B-spline derived scale-space exhibits many advantages for modeling visual mechanism with regard to the efficiency, compactness, orientation feature, and parallel structure.

Introduction Scale is a fundamental aspect of early image representation. Koenderink [1] emphasized that the problem of scale must be faced in any imaging situation. A multiscale representation is of crucial importance if one aims at describing the structure of the world. Both the psychophysical and physiological experiments have confirmed that multiscale transformed information appears in the visual cortex of mammals. This leads to the motivation for the interpretation of image structures in terms of spatial scale in computer vision. Some researchers such as Burt and Adelson [6], Koenderink [1], Marr [10], Witkin [11], Rosenfeld [2] had exposed the necessity and advantages of using operators of different sizes for extracting multiscale information in an image. For a more detailed review, see [3]. The Gaussian scale-space approach of a signal as introduced by Witkin [11], is an embedding of the original signal into a one-parameter family of derived signal constructed by convolution with Gaussian kernels of increasing width. One reason the traditional scale-space is mainly based on the Gaussian kernel is that the Gaussian function is the unique kernel which satisfy the causality property as guaranteed by the scaling theorem [14], [15], [16], [19]; it states that no new feature points are created with increasing scale. Another reason is that the response of human retina resembles a Gaussian function. Neurophysiological research by Young [22] has shown that there are receptive field profiles in the mammalian retina and visual cortex, whose measured response profiles can be well-modeled by superposition of Gaussian derivatives. Therefore, the Gaussian function is suitable for modeling human visual system. In practice, since the computational load becomes extremely heavy when the scale gets larger, many techniques are proposed for efficient implementation of scale-space filtering. Among them B-splines or binomials have been widely used to approximate the Gaussian kernel. Such examples include, Wells [5], Ferrari et al. [8], [9], Poggio et al. [20], Unser, Aldroubi and Eden [25], [26], etc. For a more compact representation, pyramid technique is another widely used representation that combine the sub-sampling operation with a smoothing step. Historically they have yielded important steps towards the scale-space theory. The low pass pyramid representation proposed by Burt [6], [7] is a famous example, which is also closely related to B-spline techniques. The general idea of representing a signal at multiple scales is not new to us. It is through wavelet theory that these early ideas have been well formulated and refined. In fact, this is largely due to the contribution of B-spline techniques. As will be shown later, the orthogonal multiresolution pyramid originally proposed by S. Mallat [40] and the biorthogonal pyramid [28], [46], [47] in wavelet theory can all be derived from B-splines [30], [33], [36]. Other types of wavelets such as the wavelets on a interval [48], the periodic wavelets [37] and the cardinal spline wavelets [29] are all related to B-splines. Motivated by these observations, the purpose of this paper is to build a more general framework of scale-space representation in the context of B-splines as an improvement of the traditional scale-space theory. By and large, the paper is divided into two parts. Firstly, we present a systematic development of the scale-space representation in the framework of B-splines. In particular, we focus on two classes of scale-space design. It is shown that if an image is represented as a B-spline efficient subdivision algorithm can be designed to give a geometric description at varying degrees of detail. The various scale-space representations are derived from B-splines in different forms. For continuous scale-space representation, a general algorithm is derived for fast and parallel implementation at rational scales. Several classical fast algorithms are shown to be special cases under some conditions. Differential operators have been used for multiscale geometric description of an image. However, it is not clear whether an image can be synthesized from these differential descriptors. Using B-spline techniques frame algorithms are designed to express the image as combinations of multiscale local partial derivatives. These operators include the gradient operator, second directional operator, Laplacian operator and oriented operators. At the same time the intrinsic relationship between wavelet theory and the traditional linear scale-space approach is exhibited. Although B-spline has been used in practice in place of Gaussian, there is little effort to consider its scale-space behavior directly. Therefore the second part of the paper is devoted to examining the properties of B-spline derived scale-space. In particular, the advantages of such a scale-space representation are highlighted. It was shown that the B-spline derived scale-space inherits most of the nice properties of the Gaussian derived scale-space. Nevertheless, the B-spline kernel outperforms the Gaussian kernel in that it can provide more meaningful, efficient and flexible description of image information for multiscale feature extraction. The organization of the paper is as follows. In Section 2, some fundamental properties of B-splines are reviewed, which also explain why a B-spline kernel is a good kernel for scale-space design. Following this section, we categorize three types of scale-space representation. In particular, we focus on the design of continuous scale-space and dyadic scale-space frame representation. The relation with compact scale-space is discussed. From the viewpoint of B-splines the evolution of wavelets from classical continuous scale-space is well understood. Moreover, the equivalence of several famous scale-space methods is explored. In Section 4 a general procedure is presented to study the edge models in the B-spline scale-space. In Section 5 some basic properties of the B-spline derived scale-space are investigated in parallel with that of the Gaussian derived scale-space. These include the the completeness property, causality property, orientation feature and so on. Conclusions are given in Section 6. Finally, Appendix A, B are given in Sections 7, 8 respectively in order to make the paper mathematically complete. 2.1 Notations and definitions We adopt the convention of [25]. Let L 2 (R) be the Hilbert space of measurable, square integrable functions on R and l 2 (Z) be the vector space of square summable sequences. We denote the central continuous B-spline of order n by fi n (x), which can be generated by repeated convolution of a B-spline of order 0, z - where the 0th-order B-spline fi 0 (x) is the pulse function with support [ ]. The Fourier transform of fi n (x) is For an integer m - 1, fi n m (x) is defined as the nth-order continuous B-spline dilated by a scale factor m, i.e., The discrete B-spline of order n at scale m is defined as z - is a normalized sampled pulse of width m. The discrete sampled B-spline b n m (k) of order n and integer coarseness m - 1 is obtained by directly sampling the nth-order continuous B-spline at the scale m: We write b n . Consequently, the frequency response of the directly sampled B-spline is an aliased version of the frequency response of the continuous B-spline since which is due to the Possion's summation formula. The discrete convolution between two sequences a and b in l 2 (Z) is the sequence b a: Under this definition, the convolution is commutative. The convolution inverse (b) \Gamma1 of a sequence b is defined by where ffi(k) is the unit impulse whose value is 1 at 0 and 0 elsewhere. The decimation operation [b] #m down-samples the sequence b by the integer factor m, i.e., Conversely, the operator [b] "m up-samples by the integer factor m, i.e., it takes a discrete signal b and expands it by padding consecutive samples: 0; otherwise: 2.2 The similarity between Gaussian and B-spline B-splines are good approximations of the Gaussian kernel which is commonly used in computer vision. This is the consequences of the central limit theorem. For a review of the B-spline window and other famous filters, see Torachi et al. [39]. In [28] Unser, Aldroubi and Eden have presented a more general proof that B-splines converge to the Gaussian function in L p (R); 8p 2 [2; +1) as the order of the spline n tends to infinity. Since the variance of the n-th order of B-spline is n+1, the approximation relation is as follows: Furthermore, by numerical computation [28] it was shown that the cubic B-spline is already near optimal in terms of time/frequency localization in the sense that its variance product is within 2% of the limit specified by the uncertainty principle. A graphical comparison between the Gaussian kernel and the cubic B-spline is given in Fig. 1. Moreover, both the physiological and biological experiments [22] have shown that the human visual system can be modeled with the Gaussian kernel. Therefore, B-splines are also suitable for modeling biological vision due to their close approximation to the Gaussian kernel. 2.3 Stable hierarchical representation of a signal by B-splines Another significant property of the B-spline of a given order n is that it is the unique compactly supported refinable spline function of order n which can provide a stable hierarchical representation of a signal at different scales. It has been proven [38] that a compactly supported spline is m-refinable and stable if and only if it is a shifted B-spline. Let h ? 0 and define the polynomial spline space S n consisting of the dilated and shifted B-splines of order n (n is odd, which we will assume throughout the paper) by Then and The embedding property (13) follows from the fact that the B-spline fi n (x) is m-refinable, i.e., it satisfies the following m-scale relation,m fi n ( x The m-refinability of the B-splines can be easily verified [29], [38]. It also establishes the intrinsic relationship between the continuous B-spline and discrete B-spline. If we take it is just the commonly used two-scale relation and B n 2 (k) is the discrete binomial. H. Olkkonen [53] has used the binomial kernels for designing multiresolution wavelet bases. From the m-scale relation (15), we can also establish the relationship between the discrete sampled B-spline and the discrete B-spline: Since B-splines provide a stable multiresolution representation of a signal at multiple scales, it is preferable to select B-splines as smoothing kernels to extract multiscale information inherent in an image. Therefore, it is not surprising that many vision models [40], [29], [30] are derived from B-splines. One can refer to [25], [26], [33], [47] for a more complete and extensive exposition of the B-spline methods. For example, its minimal support and m-refinability properties have led to fast implementation of the scale-space algorithms [27], [5], [8], [9] in computer vision. In the following section, we will classify scale-space representations into three types and show how B-splines are used as a flexible tool for designing an efficient visual model according to the above requirements. 3 Scale-space representations designed from B-splines In this section we focus on the fast implementation of continuous scale-space filtering and the design of dyadic scale-space frame representation. Their relations with the compact scale-space representation or compact wavelet models are indicated. 3.1 Implementation of continuous scale-space filtering using B-splines 3.1.1 Discrete signal approximation using B-spline bases In practice, a discrete sets of points are given. Because spline spaces S n provide close and stable approximations of L 2 , it is reasonable to parameterize the discrete signal or image using B-spline bases. We use the translated B-splines of order n 1 as bases to approximate the signal, i.e., the signal f(x) 2 L 2 (R) is projected into the spline space S n 1 1 at resolution 1. In (17) we have assumed that the sampling rate is 1 for convenience. We can call this procedure the generalized sampling of the original data, where the sampling basis function is taken as the B-spline. There are different types of approximation (see [25]). A common approach ([25], [26]) is the direct B-spline transform where the exact or reversible representation of a discrete signal f(k) in the space of B-splines is obtained by imposing the interpolation condition: 8k 2 Z; ~ f(k). Thus, the coefficient c(k) can be computed as, where (b n denotes the inverse of the discrete sampled B-spline which can be computed recursively. 3.1.2 Fast algorithm for continuous scale-space filtering at rational scales In this section we use the B-splines to derive a filter bank algorithm for fast implementation of continuous scale-space filtering. The linear scale-space representation is to make a map of a signal at multiple scales by changing the scale parameter continuously. In the language of wavelet transform the traditional scale-space approach can be regarded as a continuous wavelet transform of the signal f 2 L 2 , Z is the scaled wavelet. Because the geometric features of an image are characterized using differential descriptors, /(x) is often taken as certain derivative of a smooth kernel or has certain order of vanishing moments. Under different physical meanings various linear scale-space representations are proposed [4] where different kernels / are assumed. Here we also use the B-spline of order n 2 to approximate the wavelet /(x): In the classical scale-space theory, two frequently used multiscale edge detection filters are the famous Marr-Hildreth operator [10] and Canny operator [12], which are obtained by taking the first and second derivative of the Gaussian kernel respectively. Since B-splines are good approximations of the Gaussian kernel, we shall use the derivatives of B-splines instead. In such cases, the coefficients in (20) are the coefficients of first and second order difference operators respectively, i.e. Such spline wavelets are shown in Fig. 2. Using these spline wavelets we obtain the approximate Marr-Hildreth operator [10] and Canny operator [12] respectively. Higher order of derivatives of B-splines can detect edges with higher singularity [42] and the coefficients g are the binomial-Hermite sequences. Explicitly, in the Fourier domain the rth-order difference of the B-spline of order n can be written as which is also the rth-order derivative of the B-spline of order n + r. We remark that B-splines can also efficiently approximate other kinds of wavelets such as the generalized edge detectors in the -space representation [4] and the coefficients g can be computed numerically. Since a real number can be approximated arbitrarily close by a rational number Z, we take rational scale and derive a general filter bank implementation of the scale-space representation of (19) using the m-refinable relation (15). The cascaded implementation of (19) with / given in (20) and f approximated by (17) is, The derivation of this algorithm is given in the Appendix A. This algorithm extend that in [27]. If the scale is taken as an integer, i.e., when then the resulting formula is similar to that in [27]: c g "m 1 The implementation of the above algorithm is illustrated in the block diagram in Fig. 3. In the filter bank implementation of (23), we can interchange the order of convolution. The result is then equivalent to the discrete B-spline filtering of the difference of the discrete sampled signal (with a down-sampling) is the signal sampling followed by difference operation We note that the computational complexity is mainly due to the discrete B-spline filtering (25) which can be implemented efficiently. By (4), it turns into the cascaded convolution with the 0th- order of discrete B-spline, and can be implemented via running average sum technique. If we define such a running average operation as R then it can be realized using the following iterative strategy Therefore, starting from the initial coefficient R 0 using only the addition. Then after a down-sampling with a factor m 2 , we obtain the scale-space filtering at the rational scale m1 m2 . Suppose m 2 is fixed, as usually the case in practice, the computational cost is independent of the scale. Fig. 4 shows an example of the scale-space filtering of a simulated signal using the above algorithm. We record two important properties of this procedure. ffl Efficiency: In practice, m 2 is usually fixed. The computational complexity at each scale m 1 is O(N ). The computational complexity is largely due to the convolution with the smoothing B-spline kernel which can be reduced by running average sum technique. In contrast with the existing procedure based on direct numerical integration or FFT-based scheme, the computational complexity does not increase with the increasing number of values of the scale parameter. ffl Parallelism: The structure of the above algorithm is parallel and independent across scales. This makes it inexpensive to run on arrays of simple parallel processors. In other words this can be ideally suited for VLSI implementation. One can recall that the above subdivision algorithm is also similar to the ' a trous algorithm [49], [32] for fast computation of continuous wavelet transform except with some constraints on their filters. However, the ' a trous algorithm can only compute the wavelet transform at dyadic scales. The above algorithm can compute the wavelet transform efficiently at any scales. As will be shown in Section 3.2, if the scale is restricted to dyadic, the above algorithm is similar to the ' a trous algorithm. However, using B-splines more efficient scheme can be obtained which only need addition operation. 3.1.3 Extension to 2D images Although the above algorithm is derived in the one-dimensional case, it can be easily extended to two dimensions. The inner product B-spline fi n (x; used as a basis to parameterize the image and approximate the two-dimensional wavelet kernels. For example, we can use the tensor product B-splines to approximate the Maar-Hildreth's LoG operator [10]: where 1). Since this two-dimensional kernel is represented by the separable one-dimensional B-spline bases, by performing the above one-dimensional fast algorithm along the horizontal and the vertical orientation respectively the LoG operator can be computed efficiently. Fig. 5 shows such results for Lena image at 3 different scales. 3.2 Dyadic scale-space frame representation 3.2.1 One-dimensional signal representation by its local partial derivatives at dyadic scales The above continuous scale-space is too redundant for some applications. In addition, as stated by Witkin [11], an initial representation ought to be as compact as possible, and its elements should correspond as closely as possible to meaningful objects or events in the signal-forming process. A description that characterizes a signal by its extrema and those of its first few derivatives is a qualitative description to "sketch" a function. If we sample the scale of the above continuous scale-space as dyadic while keeping the time variable continuous, we can obtain a more compact scale-space representation. Moreover, such a representation is shift invariant and therefore is suitable for some pattern recognition applications. In particular, using B-spline techniques efficient frame algorithms can be designed to express the signal in terms of its local partial derivatives. We now show the relationship between this type of transform and the above continuous scale-space implementation. If we use the approximation b n 2 +n 1 +1 c to replace the original signal f , and the width of the B-spline m is restricted to be dyadic, say 2 m , we will get a recursive relation for the wavelet transform between the dyadic scales. In this case, the formula (23) becomes where It is easy to derive the two-scale relations for . From property (4) and the property of z-transform, is the z-transform of B n We have the following relation z \Gammaj 2 In the time domain, it becomes By this relation, we can get a fast recursive implementation of dyadic-scale space filtering, Or simply written as is the binomial kernel. The approximate Marr-Hildreth operator and Canny operator at dyadic scales can now be computed as "2 where is the first or second order of difference operator given in (21). The recursive refinement relation (36) and (37) can be rewritten in the z-transform domain as 2: (39) By requiring the reconstruction filter ~ G to satisfy the following perfect reconstruction condition we can reconstruct the signal from its multiscale partial derivatives where ~ are the time responses of ~ G k which are given explicitly in Appendix B. Since all of these filters are linear combinations of binomial and divided by 1 using Pascal triangular algorithm only addition operation and bit shift operation are needed. This is very suitable for hardware implementation. 3.2.2 Image representation by its local directional derivatives The tensor product B-spline basis is fi n (x; (y). It is interesting that we can still derive an efficient frame algorithm to characterize an image from its local differential components. A fast algorithm for the gradient case has been proposed [42] and further refined in [34]. Now we consider the case of the second directional derivative. For edge detection an approach is to detect the zero-crossings of the second directional derivative of the smoothed image f fi n along the gradient orientation @x @y We can still derive a subdivision algorithm to compute the three local partial derivative components or wavelet transforms: where the three directional wavelet components are defined in the Fourier domain as where G (1) and G (2) are the transfer functions of the first and second order difference operator. From these definitions we can obtain a recursive algorithm for the computation of the three local partial derivative components: where I (h; g) "2 j \Gamma1 represents the separable convolution of the rows and columns of the image with the one dimensional filters [h] respectively. The symbol d denotes the Dirac filter whose impulse is 1 at the origin and 0 elsewhere. We can also reconstruct the image from these dyadic wavelet transforms using the following formula 2, is the FIR of the transfer function (!). The reconstruction formula (48) follows from the following perfect reconstruct identity: G (2) (! x )U(! y G (2) (! y )U(! x ) G (1) (! y If we define the three corresponding reconstruction wavelets G (2) (! x )H 2 (! y G (2) (! y )H 2 (! x G (1) (! x ) ~ G (1) (! y it can be shown [34] that an image f(x; y) can be represented as One can also notice that in the above decomposition and reconstruction formula all the filters are binomial which require only the addition operation. For illustration Figure 6 shows the above decomposition and reconstruction results of a square image at the scales 1, 2, 4. Like the compact wavelet decomposition [40], the above algorithms also decompose an image into horizontal, vertical and diagonal components. However, this transform has explicit physical meaning and is shift-invariant. This can be useful for certain pattern recognition tasks. 3.2.3 Image representations by its isotropic and multi-orientational derivative component One can obtain a more compact isotropic wavelet representation of an image that is complete and efficient using a radial B-spline as the smoothing kernel in two dimensions. This representation is important because it indicates that an image can be recovered from its multiscale LoG-like compo- nents. The radial B-spline OE(x; y) is a non-separable function of two variables defined by its Fourier where the radius ae = min( and the wavelet is defined by One may notice that such a wavelet is isotropic and LoG-like, which resembles the human visual system. With these definitions we still have a filter bank implementation of the decomposition and reconstruction. We omit the details, and just give the decomposition formulas, r are the 2D non-separable radial filters corresponding to h and g respectively. In this decomposition, two components are obtained at each resolution. By designing the filter ~ g r from the relation (40) the corresponding 2D non-separable radial filter ~ r can be computed numerically via its Fourier transform ~ G r . Then the reconstruction formula similar to (41) can be obtained. Also, it is easy to check using the same arguments as the in 1D case that an image can be represented as [34] is the 2D reconstruction wavelet defined by - G r (ae) - OE(ae). One can build a wavelet representation having as many orientation tunings as desired by using non-separable wavelet bases. A generalized Pythagorean theorem has been proved to decompose an image into a finite number of equally spaced angles [34]: (2m\Gamma1)!!n 1). If we multiply the above isotropic wavelet (51) by the angular part H k Z, we can extract the orientational information in the dyadic scale-space tuned to n orientations [34]: Such wavelets can be called orientation tuned LoG-like filters. An image can be represented by its multiscale and multi-orientational components, is the oriented wavelet for reconstruction. Similarly, the pyramid-like filter bank implementation of such a representation can be obtained as follows, Therefore, through such an approach we can analyze the directional information of an image feature at a certain angle in dyadic scale space. In Fig. 7 we show a multiscale orientational decomposition and reconstruction, where the orientation number is chosen as 3. 3.2.4 Some comments on the application of dyadic scale-space representation The dyadic scale-space frame representation from B-spline gives rise to many applications, since it provides a invertible, translation-invariant, and pyramid-like compact representation of a signal. One example is the fingerprint based compression [41] by combining with other techniques. Many image features such as ridges, corners, blobs, junctions are usually characterized by local differential descriptors [3]. It is usually enough to consider their behaviors only at dyadic scales. The proposed algorithms provide efficient ways and are easy for hardware implementation. For example, in multi-scale shape representation usually the computation of the curvature function is treated in continuous scales [51]. In fact sometimes it is enough to consider its behavior only at the dyadic scales [50]. We have used the above algorithms to efficiently compute the geometric descriptors for multiscale shape analysis [35]. 3.3 Compact scale-space representation While the dyadic scale-space frame approach provides a more compact representation, it is still over-complete for signal representation. For image compression applications, compact representation is preferred. In order to give a complete picture we mention briefly the discrete wavelet transform. While the scale-space technique has existed for a long time, it was the orthogonal multiresolution representation proposed by S. Mallat [40] that makes the mathematical structure of the image more explicit. This is an extension or refinement of traditional scale-space theory. This approach restricts the scale to dyadic and samples the time variable. The starting point is to orthogonalize the B-spline basis, and then decompose the signal approximated at a fine scale space S n space S n by imposing the orthogonal condition The detail irregular information of the signal is contained in the subspace W n This defines an orthogonal multiscale representation. After the B-spline basis is converted into an orthogonal basis, the two-scale relation still exists which results in an efficient pyramidal algorithm. The perfect reconstruction condition (40) still exists. However, additional conditions on the filters H; G; ~ G are imposed to ensure the orthogonality. There are several ways to achieve a compact multiresolution by imposing the biorthogonal instead of the orthogonal condition (57). All these compact multiresolutions are related to B-splines. A detailed study can be found in [30], [36]. From the above analysis, it is easy to see that the dyadic scale-space frame representation lies between the continuous scale-space and the compact representation. Which kind of representation to select depends on the problem at hand. For example, in multiscale feature extraction one may compute the differential operation either at the continuous scales or only dyadic scales. Therefore, the continuous or dyadic scale-space frame representation is more useful. However, for compression applications, the compact multiresolution is the favorite. 3.4 Relations between the existing scale-space algorithms in computer vision Before the appearance of wavelets, the B-spline technique has been widely used in computer vision. Examples include Wells [5], Burt [6], [7], Ferrari et al. [8], [9]. We shall show that under certain circumstances, they are either equivalent or the special cases of the general algorithm given in Section 3. 3.4.1 Relation to Ferrari et al.'s method Ferrari et al. [8], [9] have proposed B-spline functions to realize the 2D image filtering recursively. Their idea is to use B-splines to as the filter kernel: l are the interpolation coefficients at the knots (kM; lN ), which can be computed using the usual method. Then using the properties of discrete B-spline, they derive the recursive 2D image filtering. In this way, the computational load can be greatly reduced. Therefore, this approach is a special case of our proposed algorithm for continuous scale-space filtering. 3.4.2 Relation to Wells' method In [5], Wells proposed an approach for efficient synthesis of Gaussian filters by cascaded uniform filters. It is easy to show that his method is equivalent to using cascaded 0th discrete B-spline to approximate the Gaussian kernel. By this approach, the cascaded convolution with a 0th-degree B-spline (uniform filters) can be realized by running average sum technique as discussed above. Obviously, his method is a special case of our recursive algorithm. 3.4.3 Relation to Burt's Laplacian pyramid algorithm Burt [6], [7] has introduced the following low pass filter for the generation of Gaussian or Laplacian pyramids If the parameter a is taken as a=3/8, then w(j) can be re-written as This is equivalent to the special case of in the formula (35). In this case, the filter is also equivalent to the operator used for the generation of the dual cubic spline pyramid representation as discussed in [29]. Edge patterns in the B-spline based scale-space The study of edge patterns in scale-space is very important for many applications. Much work has been done on the study of edge patterns in the Gaussian based scale-space. In this section, we want to investigate the edge behaviors in the B-spline based scale-space. It is well-known that there exists a similar "uncertainty principle" between good edge localization and noise removal. At finer scale, better localization can be achieved with the cost of noise pollution, and vice versa. Many researchers have studied the localization of operators based on Gaussian kernel in scale-space. Berzins [24] studied the accuracy of Laplacian operator, M. Shah, A. Sood and Jain [23] considered the localization of pulse and staircase edge models, I. J. Clark [21] investigated the phantom edges in a scale-space. For corner detection, A. Huertas and H. Asada [51], R. T. Chin [52] have also considered the behavior of edge models. We found that their derivations are all based on Gaussian kernel. It is necessary to study the behavior of various edge models in B-spline based scale-space which give some a priori knowledge of various patterns in an image. Here we present a more general proof with the assumption that the primitive ' in the definition dx '(x) is symmetric and with compact support [\Gammaw=2; w=2] and its derivatives have the shape as in Fig. 2. In practice, the support of the Gaussian kernel is usually truncated to a finite interval. Obviously, the truncated Gaussian and B-spline kernel are included in this assumption. We shall show that these edge models have the same behaviors as that derived from the traditional Gaussian kernel. First we adopt the following accurate definition of an edge [21]. Definition 1. A point x 0 is called an authentic edge of a signal f(x), if jW 1 f(s; x 0 )j is maxima, will be a phantom edge. We shall use W 1 f and W 2 f to denote two types of wavelet transform where the wavelets are the first and second derivatives of ' respectively. I. J. Clark [21] analyzed these two types of detection. Generally, zero-crossing detection is equivalent to the extrema detection. However, extrema detection includes both maxima and minima detection. Only the edge point detected by local maxima belongs to authentic edge and the edge point detected by local minima corresponds to phantom edge. It is shown that zero-crossing edge detection algorithms can produce edges which do not correspond to significant image intensity changes. Such edges are called phantom or spurious. Now, as an example we study the behavior of staircase edge model in scale-space. The staircase edge model can be represented as amplitudes of the edge, d is the distance between the two abrupt changes at is the step function, whose derivative in the distributional sense is ffi(x). Hence, from (19), d s (x)dx x s s i.e., the wavelet transform is just the sum of two dilated smoothing functions. At the location @ @x s s s s s Therefore, only at a small scale s ! 2d w , the location of edge at can be detected exactly. Similarly, at the location @ @x s d s s s d s i.e., the location of edge at can only be detected exactly at a small scale s ! 2d w . For large scale, it will spread like a cone in scale-space and will be influenced by another cone at As a consequence, the edge location will be mis-detected due to the superposition of two diffused cones. One may deduce that there will be another point x 0 such that @ due to different parities in the values of @ at the location of However, it is easy to find such a point corresponding to a local minimum, which means it is a phantom edge point. This cannot be distinguished from others by zero-crossing detection. Fig. 8 illustrates the behavior of this type of edge in the scale-space. That was why local maxima is preferred for edge detection in [41], [42]. Similarly, for zero-crossing detection of W 2 f(s; x), we can draw the same conclusion. In the above analysis, we only consider one type of edge model. Other types of edges such as the step, pulse, ramp, roof can be treated in a similar way [34]. Also, our derivation is based on a more general assumption on the kernels which include both the truncated Gaussian and the B-spline. 5 Discussions on the properties of the B-spline based scale-space We now discuss the advantages and properties of the B-spline based scale-space. ffl Efficiency: It is a basic requirement that any algorithm should be able to capture and process the meaningful information contained in the signal as fast as possible. Obviously, the major weakness of the traditional Gaussian based scale-space is the lack of efficient algorithms. On the contrary, B-spline techniques facilitate computational efficiency. The computational complexity is scale independent. Moreover, in contrast with the scale-space based on the Gaussian kernel, the B-spline representation of a signal is determined directly on an initial discrete set of points, avoiding problems caused by discretization in continuous scale-space. B-splines also have been used as smoothing windows for efficient computation of Gabor transform to extract frequency information [31]. ffl Parallelism: Data parallelism is common in computer vision which arises from the nature of an image. As Koendrink [1] pointed out, it seems clear enough that visual perception treats images on several levels of resolution simultaneously and that this fact must be important for the study of perception. In this paper, efficient parallel structure of an image is exhibited using B-spline techniques. It may provide a good interpretation of human visual system which can process the hierarchical information simultaneously. B-splines provide a flexible way to process the multiscale information using either coarse to fine strategy or in parallel. This is also easy for hardware implementation. ffl Completeness and invertibility: We usually use the zero-crossings or the local extrema as meaningful description of a signal. It is clearly important, therefore, to characterize in what sense the information in an image or a signal is captured by these primal sketches uniquely. For a Gaussian based scale space, the completeness property is guaranteed by the fingerprint theorem [13]: the map of the zero crossing across scales determines the signal uniquely for almost all signals in the absence of noise. Such results have theoretical interest in that they answer the question of what information is conveyed by the zero and level crossings of multiscale Gaussian filtered signals. Poggio and Yuille's proof is heavily dependent on the Gaussian kernel and they conjectured that under certain conditions Gaussian kernel is necessary for fingerprints theorem to be true. However, later Wu and Xie [17] gave a negative answer and present a more general proof, which states that fingerprint theorem holds for any symmetry kernel. Therefore, the fingerprint theorem is also true in the case of B-splines for continuous scale-space representation. Differential operators have also been widely used for multiscale geometric description of images, but it has not been clear such representations are invertible. As shown in the paper, using B- splines, efficient frame algorithms can be designed to express an image from its local derivatives at dyadic scales. ffl Compactness: For compression application, we require a representation to be as compact as possible so that an image can be represented by the corresponding primitives using less storage. In [13] Poggio and Yuille conjectured that the fingerprints are redundant and the appropriate constraints derived from the process underlying signal generation should be used to characterize how to collapse the fingerprints into a more compact representation. In the paper, the more compact dyadic scale-space representations is proposed. We can use such representation for compression applications by combining with other techniques. ffl Causality: Since edge points are important features, it is natural to require that no new features are created as the scale increases. A multiscale feature detection method that does not introduce new features as the scale increases is said to possess the property of causality. Causality is in fact equivalent to the maximum principle in the theory of parabolic differential equations [18]. The Gaussian scale-space is governed by the heat diffusion equation and therefore possesses the causality property. Such continuous causality property of the Gaussian kernel is not shared by the B-spline. However in the discrete sense, M. Aissen, I. J. Schoenberg and A. Whitney [54] proved that for a discrete scale-space kernel h, the number of local extrema or zero-crossings in f out = h f in does not exceed the number of local extrema or zero-crossings in f in if and only if its generating function H(z) =P h(n)z n is of the form where It is easy to verify that the discrete B-spline kernel in (25) satisfies such a condition. Therefore, the causality property still holds for discrete B-spline filtering in the discrete sense. The number of local extrema or zero-crossings of the derivative of the discrete signal does not increase after running average sum. This justify the use of the discrete smoothing kernel in practice. ffl Orientation: Orientation analysis is an important task in early vision and image processing, for example in texture analysis [45]. The Laplacian multiresolution [6] does not introduce any spatial orientation selectivity into the decomposition process. Daugman [45] showed that these impulse response can be approximated by Gaussian windows modulated with a wave. It is meaningful to combine both the orientation analysis and the scale feature [43]. In this paper, efficient pyramid-like algorithm is designed using B-spline technique to analyze and synthesize an image from its multi-orientational information at any number of angles in the dyadic scale-space. Note that the usual wavelet transforms can decompose an image in only three orientations. There are other advantages of the B-spline kernel. In the time-frequency analysis, the Gaussian kernel is the optimal function that minimizes the uncertainty principle. The cubic B-spline is already a good approximation to the Gaussian function [28], see also Fig. 1. As the order goes infinity, the B-splines converge to the Gaussian in both the time and frequency domain. Moreover, a B-spline resembles the response of receptive field [22] and is also suitable for modeling the human visual system. Edge detection is an ill-posed problem. From the view point of regularization theory, cubic spline is proved optimal. The connection between the regularized edge detection and the smoothing spline problem proposed by Schoenberg, Reinsh in statistics is noted by Poggio et al. [20]. It was shown that the cubic B-spline rather than the Gaussian kernel is optimal for edge detection. B-splines are the shortest basis functions that provide a stable multiresolution analysis of a signal [33]. This explains why many wavelet models of a vision [40], [28],[46], [47], [48], [37] are derived from B-splines [33], [36]. For the derivative operations, B-spline approach is very intrinsic which elucidates the relationship between derivative and difference which are usually characterized by the two-scale difference relations. B-splines play an important role to bridge the traditional scale-space theory, dyadic scale-space frame and compact multiresolution representation. 6 Conclusions This paper describes a B-spline based visual model. For a long time, the Gaussian kernel has been commonly used in computer vision. In this paper a general framework for scale-space representation using B-splines is presented. In particular, the design of two types of scale-space representations is given in detail. A fast algorithm for continuous scale-space filtering is proposed. In the case of dyadic scale some efficient frame algorithms are designed to express the image from its local differential descriptors. The intrinsic relationship with the compact wavelet models is also indicated. Several algorithms are proved to be special cases of our proposed algorithm. To our knowledge, the scale-space property based on B-splines has not been fully studied before. We examine the property of B-spline based scale-space in parallel with Gaussian kernel. Our results indicate that B-splines possess almost the same properties as the Gaussian kernel. Moreover, the B-spline kernels outperform the Gaussian in many aspects, notably, computational efficiency. Acknowledgment The authors wish to thank the referees for their comments which greatly improve the presentation of the paper. The first author wishes to express his thanks to Prof. Wu for providing the reference [17]. Appendix A: Derivation of fast implementation of continuous wavelet transform at rational scales. We use the m-scale relation (15) to derive the filter bank implementation of scale-space filtering at rational scales. From (19), (20), (17), l l Using the m-refinable relation (15), Hence, (j)B n1 (j)B n1 where the following property of B-spline is used [25]: Substituting (70) into (69) gives l l l If we take then the above formula can be written as, If we take we get an interpolation formula, and the size of the transformed data m 2 is as large as that of the original sampling signal data, Appendix B: Derivation of the reconstruction filter responses For the first order difference the perfect reconstruction condition (40) gives ~ and the corresponding FIR is j+l j+l \Gamman - l - j+l For the second order difference ~ 4z and the corresponding FIR is --R The structures of images Edge and curve detection for visual scene analysis Efficient synthesis of Gaussian filters by cascaded uniform filters The Laplacian pyramid as a compact image code Fast hierarchical correlations with Gaussian-like kernels Efficient two-dimensional filters using B-spline functions Recursive algorithms for implementing digital image filters Theory of edge detection A computational approach to edge detection Fingerprints theorems for zero-crossings Scaling theorems for zero-crossings Uniqueness of Gaussian kernel for scale-space filtering Scaling theorem for zero-crossings Reconstruction from zero-crossings in scale-space Scaling theorems for zero-crossings of bandlimited signals Computational vision and regularization theory Authenticating edges produced by zero-crossing algorithms The Gaussian derivative model for machine vision: Visual cortex simulation. Pulse and staircase edge models Accuracy of Laplacian edge detection Part II-efficient design and applications Fast implementation of continuous wavelet transforms with integer scale On the asymptotic convergence of B-spline wavelets to Gabor functions The polynomial spline pyramid A family of polynomial spline wavelet transforms Fast Gabor-like windowed Fourier and continuous wavelet transform A practical guide to the implementation of the wavelet transform Ten good reasons for using spline wavelets Multiscale curvature based shape representation using B-spline wavelets Periodic orthogonal splines and wavelets Characterization of compactly supported refinable splines Window functions represented by B-spline functions A theory for multiresolution signal decomposition: wavelet representation Characterization of signals from multiscale edges IEEE Trans. Wavelets for a vision Subband image coding using watershed and watercourse lines of the wavelet transform Complete discrete 2D Gabor transform by neural networks for image analysis and compression Biorthogonal bases of compactly supported wavelets On compactly supported spline wavelets and a duality principle Wavelets on a bounded interval The discrete wavelet transform: wedding the Multiscale corner detection by using wavelet transform The curvature primal sketch On the generating functions of totally positive sequences I --TR --CTR Anant Madabhushi , Jayaram K. Udupa , Andre Souza, Generalized scale: theory, algorithms, and application to image inhomogeneity correction, Computer Vision and Image Understanding, v.101 n.2, p.100-121, February 2006 Yu-Ping Wang , Ruibin Qu, Fast Implementation of Scale-Space by Interpolatory Subdivision Scheme, IEEE Transactions on Pattern Analysis and Machine Intelligence, v.21 n.9, p.933-939, September 1999 Yu-Ping Wang , Jie Chen , Qiang Wu , Kenneth R. Castleman, Fast frequency estimation by zero crossings of differential spline wavelet transform, EURASIP Journal on Applied Signal Processing, v.2005 n.1, p.1251-1260, 1 January 2005 Huamin Feng , Wei Fang , Sen Liu , Yong Fang, A new general framework for shot boundary detection and key-frame extraction, Proceedings of the 7th ACM SIGMM international workshop on Multimedia information retrieval, November 10-11, 2005, Hilton, Singapore Luca Lombardi , Alfredo Petrosino, Distributed recursive learning for shape recognition through multiscale trees, Image and Vision Computing, v.25 n.2, p.240-247, February, 2007

b-spline;image modeling;scaling theorem;wavelet;scale-space;fingerprint theorem

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Efficient Error-Correcting Viterbi Parsing.

AbstractThe problem of Error-Correcting Parsing (ECP) using an insertion-deletion-substitution error model and a Finite State Machine is examined. The Viterbi algorithm can be straightforwardly extended to perform ECP, though the resulting computational complexity can become prohibitive for many applications. We propose three approaches in order to achieve an efficient implementation of Viterbi-like ECP which are compatible with Beam Search acceleration techniques. Language processing and shape recognition experiments which assess the performance of the proposed algorithms are presented.

Introduction The problem of Error-Correcting Parsing (ECP) is fundamental in Syntactic Pattern Recognition (SPR) [11, 13], where data is generally distorted or noisy. It also arises in many other areas such as Language Modeling [23, 6], Speech Processing [7, 22], OCR [18], Grammatical Inference [20], Coding Theory [8, 15] and Sequence Comparison [21]. As in many other problems arising in several research areas, ECP is related to finding the best path through a trellis, a problem that is solved by the Viterbi algorithm (VA) [10] as is well-known. McEliece [19] makes a good description of the operation and complexity of the VA with application to decoding linear block codes. On the other hand, Wiberg et al [24] use an alternative decoding scheme based on Tanner graphs rather than trellises. Nevertheless, only the work described in [8, 15] seems to specifically deal with the ECP problem as stated in Sect. 2. Henceforth we shall be concerned with the application of ECP in SPR that is, it will be assumed that both a stochastic Finite State Model (FSM) and an (stochas- tic) error model accounting for insertions, substitutions and deletions of symbols are given 1 . These symbols belong to some alphabet \Sigma, which stands for the set of primitives or features that characterise a given pattern we aim to recognise. Therefore, objects are represented by strings of symbols belonging to \Sigma. On the other hand, the FSM accounts for the (generally infinite) set of different strings corresponding to the several ways a given object can be represented, while the error model accounts for the typical variations that pattern strings tend to exhibit with regard to their "standard" forms as represented by the FSM. If no error model is given, the recognition problem amounts to a simple problem of Finite-State parsing. Given an input string of symbols, we have to compute the probability that this string belongs to the language generated by the FSM. From this point of view, we are only interested in the maximum likelihood derivation ("Viterbi-like") of the string, instead of the sum of the likelihoods of every derivation ("Forward-like"). In a multiclass situation, where one FSM is provided for each class, these probabilities can be used for recognition using the maximum likelihood classification rule: the string is classified into the class represented by the FSM whose probability of generating the string is maximum. However, in many cases, test strings cannot be exactly parsed through any of these FSMs, leading to zero probabilities for all the classes. This can often be solved through ECP. If the FSM is deterministic, (non error-correcting) parsing is trivial. Otherwise, the VA can be used. If an error model is provided, the same Viterbi framework can be adopted for ECP, but at the expense of a higher computational cost. Unfor- 1 Note that this work is also related to the problem of (approximately) matching regular expres- sions, since they are equivalent to FSMs. See [2] for an introduction to this problem. tunately, this higher cost can become prohibitive in many applications of interest. The computational problem of ECP is outlined in the next section. Solutions to this problem are proposed in Sects. 3, 4 and 5. Sect. 6 describes the adaptation of the well-known Beam Search technique [17] to further accelerate the parsing process. Sect. 7 presents the experiments that have been carried out to test the performance of the distinct approaches. 2 The Computational Problem of ECP In general, the problem of Finite-State parsing with no error correction can be formulated as a search for the "minimum cost" path 2 through a trellis diagram associated to the FSM and the given input string, x. This trellis is a directed acyclic multistage graph, where each node q j k corresponds to a state q j in a "stage" k. The stage k is associated with a symbol, x k , of the string to be parsed and every edge of the trellis, t stands for a transition between the state q i in stage k and the state q j in stage k Thanks to the acyclic nature of this graph, Dynamic Programming (DP) can be used to solve the search problem, leading to the well-known Viterbi algorithm. The trellis diagram can be straightforwardly extended to parse errors produced by changing one symbol for another symbol and errors produced by inserting a symbol before or after each symbol in the original string. In this way, taking only substitution and insertion errors into account, efficient ECP can be implemented because such an extended trellis diagram still has the shape of a directed acyclic multistage graph (Fig. 1 (a) (b)). Unfortunately, the extension of this trellis diagram to also parse errors produced by deletion of one or more (consecutive) symbol(s) in the original string results in a kind of graph with edges between nodes belonging to the same stage k (Fig. 1 (c)). Nevertheless, if the FSM has no cycles, the resulting graph is still acyclic and DP negative log probabilities and sums rather than products are used to avoid underflows. (d) (c) (b) (a) K K K K Figure 1: Trellis a) Substitution and proper FSM transitions b) Insertion transitions c) Deletion transitions in an acyclic FSM d) Deletion transitions in a cyclic FSM. Each edge is actually labelled with a symbol of \Sigma can still be applied, leading to an efficient algorithm that can be implemented as a simple extension of VA [3]. However, if cycles exist in the FSM, then DP can no longer be directly used and the problem becomes one of finding a minimum cost path through a general directed cyclic graph (Fig. 1 (d)). As noted in [15], we can still take advantage of the fact that most of the edges of this kind of graphs have a left-to-right structure and consider each column as a separate stage like in the VA. 3 Solving the Problem by Score Ordering Bouloutas et al [8] formulate an interesting recurrence relation to solve the problem stated in Sect. 2. In our notation it is as follows: l ) in stage k 8l in stage k (1) ) is the cost of the minimum cost path from any of the initial states to state q i in stage k; is the inverse of the transition function, ffi, of the FSM; k+1 ) is the cost of the minimum cost transition from state q i in stage k to state q l in stage k k+1 ) is the cost of the minimum cost path from state q l to state q j , both of which are in stage k + 1. Its correctness lies in i , since for each pair of states of the FSM its evaluation yields the cost of the minimum cost deletion path between them. Fig. 2 shows the algorithm, called EV1, which we developed from (1), with Q being the set of states of the FSM. Lines 1-3 will be referred to as init-block, lines 12-16 as ins-subs-block and line as ret-block in the remainder of the paper. 1. for each j 2 Q do 2. if j is an initial state then C(q j InitialCost else C(q j 3. endfor 4. for to jxj do 5. do 7. 8. for each 9. C(q j 10. endfor 11. endwhile 12. for each i 2 Q do 13. for each l 2 ffi(q i ) do 14. C(q l C(q l 15. endfor 16. endfor 17. endfor 18. return argmin for each final state i Figure 2: Algorithmic scheme for EV1 and EV1PQ Given that there are no transitions with a negative cost, Dijkstra's strategy is followed in order to compute i (lines 6-11 in Fig. 2). All transitions from a state to itself can be discarded to perform this computation. The state q l whose score in a given parsing stage is minimum is chosen and the score of each j 2 ffi(q l ) is updated. Again, the minimum-score state is chosen and the scores of its direct successors are updated, and so on, until there are no states whose score could be updated. An input string x can be parsed in \Theta(jxj \Delta using EV1 if no care about the implementation of Q 0 is taken. This algorithm can be further improved by using priority queues [1] in the implementation of Q 0 . In this implementation, the scores (therefore the positions) of the states in the heap need to be dynamically changed. This can be done by simply storing the pointer to each state in the heap in order to perform a heapify operation from the position of the state whose score has changed. The worst-case time complexity of the loop in lines 6-11 of Fig. 2 is, in this case, O(jQj given that at most jQj \Delta B operations will be performed in the heap [1], and with B being the "Branching Factor" or maximum number of transitions associated to each state in Q. Since B ! jQj in many cases, the worst-case time complexity of this version of EV1 (called EV1PQ throughout the paper) to parse an input string x is O(jxj log jQj). Note that if the FSM is a fully-connected graph, then and the performance of EV1PQ can be worse than that of EV1. 4 Solving the Problem Iteratively Another approach for coping with the deletion problem consists in performing consecutive iterations to compute the minimum cost path to each state in any parsing stage k (line 9 in Fig. 2), using only deletion transitions (Fig. 1 (c) (d)). If this iterative procedure is performed until no score updating is produced then the properness of the overall computation is guaranteed. This idea was independently proposed in [20] and [15], though the paper by Hart and Bouloutas is more comprehensive. Their work deals with many kinds of error rules and efficiently copes with their associated computational problems. The resulting algorithm, called EV2, is shown in Fig. 3. Let T be the number of iterations to be done at lines 3-9 in Fig. 3. The time complexity of these lines is O(T \Delta jQj \Delta B). T would be 1 if no deletion transition changed the score of any state of the FSM [5]. If, at least, one deletion transition per state changed the score of any other state through consecutive iterations then T would be jQj can be produced if the states are traversed in reverse topological or score order [5]). EV2 can thus parse an input string x in O(jxj \Delta jQj \Delta B) 1. init-block 2. for to jxj do 3. repeat 4. for each l 2 Q do 5. for each j 2 ffi(q l ) do 7. endfor 8. endfor 9. until C(q j k ) has not been changed for any j 2 Q 10. ins-subs-block 11. endfor 12. ret-block Figure 3: Algorithmic scheme for EV2 and O(jxj \Delta jQj 2 \Delta B) in the best and worst cases, respectively. Its performance in the average case strongly depends on following some order of the states of the FSM as closely as possible when parsing deletion transitions and on the number of "effective" deletion transitions [5]. 5 Solving the Problem by Depth-First Ordering Here, we propose an algorithm based on a recurrence relation which extends previous ideas of [20] for ECP with acyclic FSMs to general FSMs: l ) in stage k (2) are the same as in equation (1); is a generalisation of function which returns, for a given state q, the set of states that are "topological predecessors" of q in the FSM, that is is an initial state and W T (q l k+1 ) is the minimum cost path from state q l to state q j , both of which are in stage k + 1, which only includes states that are topological predecessors of state q j . The computation of W T can be performed by following a "topological order" of the states of the FSM when parsing deletion transitions. If the FSM has cycles, a depth-first "topological sort" of its states can be computed by detecting the back-edges [3] (i.e., transitions which produce cycles in the FSM). This leads to a fixed order for the traversal of the list of states of the FSM during the parsing process. Backtracking becomes necessary to ensure the correctness of the overall computation only if some back-edge coming from a state q i to another state q j updates the cost of q j . This is a solution for the problem stated in Sect. 2, but, unfortunately, it is not directly compatible with Beam Search (BS) techniques. When using BS the list of states to be traversed can be different in almost every parsing stage; therefore large computational overheads could be introduced during the parsing process if we had to compute a depth-first sort of the list of "visited states" in each parsing stage. This can be avoided by depth-first sorting the states as they are visited, using bucketsort (binsort) techniques [1]. We only need to use an adequate ordering key. Our proposal is to compute and store the ordering key \Psi i 8i 2 Q as shown in Fig. 4. 1. for each i 2 Q do 2. #back-edges coming to i 3. endfor 4. for each i such that ae do 5. for each such that the edge (q is not a back-edge do 7. ae 8. endfor 9. endfor Figure 4: Computation of \Psi i 8i 2 Q. Both \Psi and ae have been implemented as arrays It can be easily shown that the relation on the set \Psi i 8i 2 Q is a partial order [1]. Therefore, no pair of states q i and q j such that \Psi exists having a transition (path) from q i to q j and vice versa. A permutation that maps Q into a nondecreasing sequence \Pi(Q) can be found using \Psi by means of bucketsort. The number of buckets to be used is max . The depth-first traversal of the FSM, along with the computation of \Psi i 8i 2 Q, can be performed in a preprocessing stage taking only O(jQj \Delta B) computing steps [4, 5]. Once \Pi(Q) has been found, the only thing to worry about in a given parsing stage is to find out when a back-edge is parsed. If a back-edge is being parsed. Backtracking is performed only if the score of q j is changed. Two algorithms based on equation (2), EV3 and EV3.V2, have been developed [5]. In both algorithms \Pi(Q) has been implemented as a hash table. Fig. 5 shows EV3 algorithm. EV3.V2 is similar but it uses only the list \Pi(Q), performing the parsing of insertion, substitution and deletion errors at once, while EV3 uses both the (unsorted) list Q (for BS purposes, see next section) in the parsing of insertions and substitutions and the list \Pi(Q) in the parsing of deletions. 1. init-block 2. 3. for to jxj do 4. for each l 2 \Pi(Q) do 5. for each j 2 ffi(q l ) do 7. if j 62 \Pi(Q) then Add j to bucket \Psi j in \Pi(Q) endif 8. if C(q j k ) has been changed and \Psi 9. then backtrack to bucket \Psi j in \Pi(Q) endif 10. endfor 11. endfor 12. for each i 2 Q do 13. for each l 2 14. C(q l C(q l 15. if l 62 \Pi(Q) then Add l to bucket \Psi l in \Pi(Q) endif 16. endfor 17. endfor 18. endfor 19. ret-block Figure 5: Algorithmic scheme for EV3 The time complexity of lines 4-11 in Fig. 5 is the time of finding \Pi(Q) times the maximum branching factor (B). \Pi(Q) is found using bucketsort. The complexity of bucketsort to sort n elements is O(n +m), with m being the number of buckets. In this case \Psi i is clearly bounded by jQj (see Fig. 4). In the best case (no back-edge requires a backtracking recomputation) the resulting time complexity is, therefore, O(jQj \Delta B). In the worst case there will be a linear-on- jQj number of back-edges requiring a backtracking recomputation of all the paths already computed 3 [20], leading to O(jQj 2 \Delta B) time complexity. EV3 and EV3.V2 can parse an input string x in O(jxj \Delta jQj \Delta B) and O(jxj \Delta jQj 2 \Delta B) in the best and worst cases, respectively. The performance on the average case not only depends on the structure of FSMs but also on the number of back-edges that require a backtracking computation from the state that is reached by them. A theoretical formulation of this average cost is very difficult and it requires assumptions on probabilistic distributions over the space of possible FSMs, which is not always feasible. 6 Beam Search Beam Search [17] (BS) is a classical acceleration technique for VA. This search technique often yields approximately optimal (or even optimal) solutions while drastically cutting down the search space [22, 5]. In this respect, BS is comparable to other "clever" strategies based on A search, such as that proposed in [14]. The idea is to keep only a "beam" of "promising" paths at each stage of the trellis. The beam width is a constant value to be empirically tuned to achieve an adequate tradeoff between the efficiency and the accuracy of the search. The lower this parameter, the lower the accuracy and the lower the computing time, and vice versa. The implementation of BS consists in keeping only the paths ("visited states") with a score which is lower than a given bound. For the sake of efficiency, this bound is implemented by adding the currently found lowest score to the beam width [22]. Q is implemented as a double linked list, leaving the first place for this lowest-score state, to avoid overhead. This strategy generally results in no significant differences 3 This upper bound is quite pessimistic since it assumes that after the depth-search ordering, the resulting number of back-edges that change the score of some state is linear on jQj. with regard to a strict implementation of BS. The extension of the algorithms EV1, EV1PQ and EV2 to perform BS is straightforward [4, 5]. In the case of EV3 and EV3.V2 the extension is easy thanks to the ordering key \Psi, which allows for building the list \Pi(Q) as the states are visited (see Sect. 5). EV1, EV1PQ and EV2 use the linked list, but slight differences in the number of visited states can exist due to the fact that EV1 and EV1PQ parse deletion transitions in score order. EV3 also uses the linked list to be able to tightly follow this BS strategy. Again, slight differences can exist, but only due to the fact that EV3 parses only the deletions in "topological" order while EV3.V2 parses all transitions following this order. This is a problem for the computation of the first bound value to be used in the next parsing stage k since it is likely that the first state in \Pi(Q) is not the minimum-score one. EV3.V2 overcomes this problem by computing an approximate bound value as the minimum cost edge of the lowest-score state found in parsing stage k plus the beam width. In practice, the differences in the number of visited states among all algorithms prove to be negligible [5]. 7 Experiments and Results Two series of experiments were carried out in order to assess both the effectiveness (parsing results) of ECP and (mainly) the efficiency (speed) of the algorithms previously discussed. In the first case, ECP was used to "clean" artificially distorted sentences from a Language Learning task called "Miniature Language Acquisition" In the second case, ECP was applied to recognise planar shapes (hand- written digits), coded by chain-coding the contours of the corresponding images [18]. In both cases, the required stochastic FSMs were automatically learned by means of the k-TSI Grammatical Inference algorithm proposed in [12]. This algorithm infers a (stochastic) FSM that accepts the smallest k-Testable Language in the Strict Sense (k-TS language) that contains the training sentences. Stochastic k-TS 4 This consists in a pseudo-natural language for describing simple visual scenes, with a vocabulary of 26 words. languages are equivalent to the languages modeled by the well-known N-GRAMS, with Increasing values of k from 2 to 10 (2 to 7) were used in the first (second) series of experiments. This yielded increasing size FSMs as required for studying the computational behaviour of the different algorithms. In all cases, only roughly hand-tuned error-model parameters were used. All the experiments were carried out on an HP9000 Unix Workstation (Model 735) performing 121 MIPS. 7.1 Language Processing Experiments A set of nine stochastic FSMs, ranging from 26 to 71; 538 states, which were automatically learned from 50; 000 (clean) sentences of the MLA task [9] were used in these experiments. The test set consisted of 1; 000 sentences different from those used in training. It was distorted using a conventional distortion model [16] in order to simulate the kinds of errors typically faced in speech processing tasks. This generally resulted in sentences which no longer belonged to the languages of the learned FSMs. Three different percentages of global distortion, which were evenly distributed among insertion, deletion and substitution parameters, were used: 5% and 10%. The effectiveness of ECP is summarised in Table 1. The quality of parsing was measured in terms of both word error rate (WER) 5 between the original (undis- torted) test sentences and those obtained by ECP (from the corresponding distorted sentences), and test-set Perplexity (PP) 6 . These results were obtained without BS and were identical for all the algorithms. For adequately learned k-TS models the distortion (WER) of test sentences was reduced by a factor ranging from 2 to 3; the best results obtained by the 6-TS model with 3; 231 states. Perplexity figures closely follow the tendency of WER. Values of k greater than 6 tended to degrade the results, due to the lack of generalisation usually exhibited 5 minimum number of insertions, substitutions and deletions. 6 2 to the power of the Cross-Entropy, which is the sum of the maximum log-likelihood score for each input distorted test sentence divided by the overall number of parsed words [23]. by k-TS models as k is increased beyond a certain value (6 in this case). This is explicitly assessed by the column labelled N in Table 1, which shows the number of original (undistorted) test sentences that would not have been accepted by the corresponding k-TS models without ECP 7 . Table 1: Parsing results in terms of Word Error Rate (WER-%) and test-set Perplexity (PP) for each FSM without BS in the LP experiments Distortion 1% Distortion 5% Distortion 10% jQj (value of 26 (2) 0 0.43 4.36 2.19 5.55 4.80 7.28 71; 538 (10) 475 3.73 3.48 4.87 4.15 6.32 5.29 Table 2: WERs and PPs for each Beam width using the FSM with 3; 231 states in the LP experiments Distortion WER PP WER PP WER PP 5% 3.46 4.23 1.70 3.87 1.70 3.87 10% 6.08 5.71 3.59 5.17 3.32 5.15 Table 2 shows the effect of using BS in the ECP process for the 6-TS model (3; 231 states) using EV1. Only negligible differences were observed for the other algorithms, due to the distinct ways BS is implemented [5]. Four increasing values of beam width (ff) were tested: 5, 10, 20 and 40 (1 means no BS). Setting ff to 20 provided results identical to those achieved by a full search (see also Table 1), while a further decreasing of ff gracefully degrades the results. Fig. 6 shows the relative efficiency of the different algorithms for increasing FSM 7 It is notable that even with almost 50% undistorted test sentences rejected without ECP, a noticeable distortion reduction is still achieved for 10% and 5%-distorted sentences. Therefore, proves useful not only in dealing with imperfect input, but also in improving the effectiveness of imperfect FSMs. Computation time (in centiseconds) Number of states in FSM Figure Average computing times (in centiseconds) of the different algorithms measured without BS for 10% of distortion in the LP experiments sizes without BS. Only results for 10% of distortion are reported; the results were similar for 5% and 1%. The variance observed in these results was negligible. More specifically, the standard deviation of the computing time per symbol parsed was never greater than the 6:7% of the average computing time per symbol parsed. This means that, with a probability close to 1, the real computing times will match their corresponding expected values. A dramatically higher computational demand of EV1 is clear in this figure. The use of priority queues (EV1PQ) to compute i contributes to alleviating the computational cost, though the resulting jQj: log jQj time complexity is still exceedingly high. The two implementations of EV3 show a much better (linear-on-jQj) performance 8 . EV2 also shows linear time complexity but with a slope which is larger than that of EV3. This is due to the number of iterations required for the parsing of deletion transitions [5]. Figs. 7 and 8 show the impact of BS in the performance of the different algorithms. Fig. 7 shows the effect of increasing the distortion rate for a fixed ff 9 (20; i.e., results 8 The observed computing times of the preprocessing stage for EV3 and EV3.V2 (see Sect. 5) were negligible, ranging from less than a centisecond (for the smallest FSM) to 36 centiseconds (for the largest FSM). 9 Differences between scores through consecutive parsing stages tend to level as larger errors are produced. Therefore, an increase in the distortion produces an increase in the rate of visited states 1% dist.0,2 0,4 0,6 Computation time (in centiseconds) Number of states in FSM0,51,52,53,54,55,5 Computation time (in centiseconds) Number of states in FSM 5% dist.0,51,52,53,54,55,50 10000 20000 30000 40000 50000 60000 70000 80000 Computation time (in centiseconds) Number of states in FSM 10% dist. Figure 7: Average computing times observed for each distortion rate and in the LP experiments identical to full search). With 1% of distortion, the number of visited states is very small and all the algorithms perform fairly well. With 5% and 10% of distortion, EV1 is again highly cost-demanding. EV1PQ also tends to have higher computational costs. EV2 and EV3 show the best performance and, by also using BS, EV3 gets the best results for the larger FSMs. Finally, Fig. 8 shows the effect of the different ff for the highest distortion rate (10%). Significant differences exist for ff 20. It should be taken into account that when ff is smaller than 20, only suboptimal results are produced (see Tables 1 and 2). The best results are systematically achieved by EV3 and EV2, although EV3 clearly outperforms EV2 when the number of visited states is higher than a certain bound. See [5] for further details on this experimental study. 7.2 OCR Experiments In these experiments, 2; 400 images of handwritten digits were used. Strings representing these images were obtained by following their outer contour using a chain-code of eight directions [18]. The resulting string corpus was randomly split into (for a given ff). 0,4 0,6 Computation time (in centiseconds) Number of states in FSM 0,4 0,6 Computation time (in centiseconds) Number of states in FSM Computation time (in centiseconds) Number of states in FSM Computation time (in centiseconds) Number of states in FSM Figure 8: Average computing times (in centiseconds) observed for 10% of distortion and Beam width values 5, 10, 20 and 40 in the LP experiments disjoint training and test sets of the same size. The training set was used to automatically learn six different stochastic k-TS FSMs per digit using values of k from 2 to 7. For each value of k, the 10 k-TS models learned for each of the 10 classes (digits) were merged into a whole global model, with class-dependent labelled final states. This resulted in six stochastic FSMs of increasing size, ranging from 80 to states. Testing was carried out using the maximum likelihood classification rule as mentioned in Sect. 1. Table 3 shows the overall recognition rates achieved with ECP without BS. The FSM with 18; 416 states gets the best results; further decreasing of the value of k tended to degrade the effectiveness. Informal tests show that recognition rates without ECP drop below 50% for all the models. Table 3: Recognition rate (%) achieved in the OCR experiments for each FSM jQj (value of Recog. rate 75.8 90.8 95.1 97.2 97.2 97.5 Computation time (in centiseconds) Number of states in FSM Figure 9: Average computing times (in centiseconds) of the different algorithms measured without BS in the OCR experiments Fig. 9 shows the observed average parsing times for each ECP algorithm without BS. The observed variance was also negligible (the observed standard deviation of the computing time per symbol parsed was never greater than the 7:2% of the average computing time per symbol parsed). Again, the performance achieved by much better than that of the other algorithms. The performance of EV2 was, in this case, worse than the one observed in the previous set of experiments (in even EV1PQ outperformed EV2). This is because EV2 significantly depends on specific conditions of data (see [5]). The impact of BS was studied for the model supplying the best recognition results; namely the 7-TS model with 18; 416 states (see Table 3). The results are shown in Table 4. In this case, none of the values of ff tested yielded identical recognition rate as that achieved by full search (97:5%). However, a width of 30 appears to be a good tradeoff between efficiency and recognition rate. The fastest performance using BS was again achieved by EV2 and EV3. The differences among the different algorithms were more significant for the widest beam (30). In this case, EV3 was about 1.5 times faster than EV2 and 1.7 times faster than EV1PQ. 10 The observed computing times of the preprocessing stage for EV3 and EV3.V2 were again negligible, ranging from less than a centisecond to 7 centiseconds. Table 4: Impact of using BS in recognition rate (%) and computing time (centiseconds) in the OCR experiments ff Recog. rate EV3 EV3.V2 EV2 EV1PQ EV1 Concluding Remarks Several techniques have been proposed for a cost-efficient implementation of Finite-State Correcting Viterbi Parsing. This is a key process in many applications in areas such as Syntactic Pattern Recognition, Language Processing, Grammatical Inference, Coding Theory, etc. A significant improvement in parsing speed with regard to previous approaches can be achieved by the EV3 algorithm which is proposed here. Furthermore, a dramatic acceleration can be achieved by applying suboptimal Beam Search strategies to the proposed algorithms. All the algorithms developed allow for integration with this search strategy, although minor differences in the rate of visited states lead to small differences in performance. In this case, EV3 has also exhibited better behaviour than all the other algorithms. 9 Acknowledgements The authors wish to thank the anonymous reviewers for their careful reading and valuable comments. Work partially funded by the European Union and the Spanish CICYT under contracts IT-LTR-OS-30268 and TIC97-0745-C01/02. --R The Design and Analysis of Computer Algorithms. "Algorithms for Finding Patterns in Strings" "Fast Viterbi decoding with Error Correction" "Two Different Approaches for Cost-efficient Viterbi Parsing with Error Correction" "Different Approaches for Efficient Error-Correcting Viterbi Parsing: An Experimental Comparison" "Simplifying Language through Error-Correcting Decoding" "Decoding for Channels with Insertions, Deletions and Substitutions with Applications to Speech Recognition" "Two Extensions of the Viterbi Algorithm" "Miniature Language Acquisition: A touchstone for cognitive science" "The Viterbi algorithm" "Inference of k-testable languages in the strict sense and application to An Introduction. "Efficient Priority-First Search Maximum-Likelihood Soft-Decision Decoding of Linear Block Codes" "Correcting Dependent Errors in Sequences Generated by Finite-State Processes" "Evaluating the performance of connected-word speech recognition systems" "The Harpy Speech Recognition System" "A Comparison of Syntactic and Statistical Techniques for Off-Line OCR" "On the BCJR Trellis for Linear Block Codes" Time Warps "Fast and Accurate Speaker Independent Speech Recognition using structural models learnt by the ECGI Algorithm" "Codes and Iterative Decoding on General Graphs" --TR --CTR Christoph Ringlstetter , Klaus U. Schulz , Stoyan Mihov, Orthographic Errors in Web Pages: Toward Cleaner Web Corpora, Computational Linguistics, v.32 n.3, p.295-340, September 2006 Juan-Carlos Amengual , Alberto Sanchis , Enrique Vidal , Jos-Miguel Bened, Language Simplification through Error-Correcting and Grammatical Inference Techniques, Machine Learning, v.44 n.1-2, p.143-159, July-August 2001 Francisco Casacuberta , Enrique Vidal, Learning finite-state models for machine translation, Machine Learning, v.66 n.1, p.69-91, January 2007

depth-first topological sort;shape recognition;viterbi algorithm;priority queues;error-correcting parsing;beam search;sequence alignment;language processing;bucketsort

290216

Simple Fast Parallel Hashing by Oblivious Execution.

A hash table is a representation of a set in a linear size data structure that supports constant-time membership queries. We show how to construct a hash table for any given set of n keys in O(lg lg n) parallel time with high probability, using n processors on a weak version of a concurrent-read concurrent-write parallel random access machine (crcw pram). Our algorithm uses a novel approach of hashing by "oblivious execution" based on probabilistic analysis. The algorithm is simple and has the following structure: Partition the input set into buckets by a random polynomial of constant degree. For t:= 1 to O(lg lg n) do Allocate Mt memory blocks, each of size Kt. Let each bucket select a block at random, and try to injectively map its keys into the block using a random linear function. Buckets that fail carry on to the next iteration. The crux of the algorithm is a careful a priori selection of the parameters Mt and Kt. The algorithm uses only O(lg lg n) random words and can be implemented in a work-efficient manner.

Introduction Let S be a set of n keys drawn from a finite universe U . The hashing problem is to construct a with the following attributes: Injectiveness: no two keys in S are mapped by H to the same value, Space efficiency: both m and the space required to represent H are O(n), and Time efficiency: for every x 2 U , H(x) can be evaluated in O(1) time by a single processor. Such a function induces a linear space data structure, a perfect hash table, for representing S. This data structure supports membership queries in O(1) time. This paper presents a simple, fast and efficient parallel algorithm for the hashing problem. Using n processors, the running time of the algorithm is O(lg lg n) with overwhelming probability, and it is superior to previously known algorithms in several respects. Computational models As a model of computation we use the concurrent-read concurrent- parallel random access machine (crcw pram) family (see, e.g., [35]). The members of this family differ by the outcome of the event where more than one processor attempts to write simultaneously into the same shared memory location. The main sub-models of crcw pram in descending order of power are: the Priority ([29]) in which the lowest-numbered processor succeeds; the Arbitrary ([42]) in which one of the processors succeeds, and it is not known in advance which one; the Collision + ([9]) in which if different values are attempted to be written, a special collision symbol is written in the cell; the Collision ([15]) in which a special collision symbol is written in the cell; the Tolerant ([32]) in which the contents of that cell do not change; and finally, the less standard Robust ([7, 34]) in which if two or more processors attempt to write into the same cell in a given step, then, after this attempt, the cell can obtain any value. 1.1 Previous Work Hash tables are fundamental data structures with numerous applications in computer science. They were extensively studied in the literature; see, e.g., [37, 40] for a survey or [41] for a more recent one. Of particular interest are perfect hash tables, in which every membership query is guaranteed to be completed in constant time in the worst case. Perfect hash tables are perhaps even more significant in the parallel context, since the time for executing a batch of queries in parallel is determined by the slowest query. Fredman, Koml'os, and Szemer'edi [16] were the first to solve the hashing problem in expected linear time for any universe size and any input set. Their scheme builds a 2-level hash function: a level-1 function splits S into subsets ("buckets") whose sizes are distributed in a favorable manner. Then, an injective level-2 hash function is built for each subset by allocating a private memory block of an appropriate size. This 2-level scheme formed a basis for algorithms for a dynamic version of the hashing problem, also called the dictionary problem, in which insertions and deletions may change S dynamically. Such algorithms were given by Dietzfelbinger, Karlin, Mehlhorn, Meyer auf der Heide, Rohnert and Tarjan [12], Dietzfelbinger and Meyer auf der Heide [14], and by Dietzfelbinger, Gil, Matias and Pippenger [11]. In the parallel setting, Dietzfelbinger and Meyer auf der Heide [13] presented an algorithm for the dictionary problem. for each fixed ffl - 0, n arbitrary dictionary instructions (insert, delete, or lookup), can be executed in O(n ffl ) expected time on a a n 1\Gammaffl -processor Priority crcw. Matias and Vishkin [39] presented an algorithm for the hashing problem that runs in O(lg n) expected time using O(n= lg n) processors on an Arbitrary crcw. This was the fastest parallel hashing algorithm previous to our work. It is based on the 2-level scheme and makes extensive use of counting and sorting procedures. The only known lower bounds for parallel hashing were given by Gil, Meyer auf der Heide and Wigderson [27]. In their (rather general) model of computation, the required number of parallel steps is \Omega\Gamma/1 n). They also showed that in a more restricted model, where at most one processor may simultaneously work on a key, parallel hashing time is\Omega\Gamma/3 lg n). They also gave an algorithm which yields a matching upper bound if only function applications are charged and all other operations (e.g., counting and sorting) are free. Our algorithm falls within the realm of the above mentioned restricted model and matches the \Omega\Gammae/ lg n) lower bound while charging for all operations on the concrete pram model. 1.2 Results Our main result is that a linear static hash table can be constructed in O(lg lg n) time with high probability and O(n) space, using n processors on a crcw pram. Our algorithm has the following properties: Time optimality It is the best possible result that does not use processor reallocations, as shown in [27]. Optimal speed-up can be achieved with a small penalty in execution time. It is a significant improvement over the O(lg n) time algorithm of [39]. Reliability Time bound O(lg lg n) is obeyed with high probability; in contrast, the time bound of the algorithm in [39] is guaranteed only with constant probability. Simplicity It is arguably simpler than any other hashing algorithm previously published. (Never- theless, the analysis is quite involved due to tight tradeoffs between the probabilities of conflicting events.) Reduced randomness It is adapted to consume only O(lg lg n) random words, compared to \Omega\Gamma n) random words that were previously used. Work optimality A work optimal implementation is presented, in which the time-processor product is O(n) and the running time is increased by a factor of O(lg n); it also requires only O(lg lg n) random words. Computational model If we allow lookup time to be O(lg lg n) as well, then our algorithm can be implemented on the Robust crcw model. Our results can be summarized in the following theorem. Theorem 1 Given a set of n keys drawn from a universe U , the hashing problem can be solved using O(n) space: (i) in O(lg lg n) time with high probability, using n processors, or (ii) in O(lg lg n lg n) time and O(n) operations with high probability. The algorithms run on a crcw pram where no reallocation of processors to keys is employed, and use O(lg lg bits. The previous algorithms implementing the 2-level scheme, either sequentially or in parallel, are based on grouping the keys according to the buckets to which they belong, and require learning the size of each bucket. Each bucket is then allocated a private memory block whose size is dependent on the bucket size. This approach relies on techniques related to sorting and counting, which require \Omega\Gammaeq n= lg lg n) time to be solved by polynomial number of processors, as implied by the lower bound of Beame and Hastad [4]. This lower bound holds even for randomized algorithms. (More recent results have found other, more involved, ways to circumvent these barriers; cf. [38, 3, 26, 30].) We circumvent the obstacle of learning buckets sizes for the purpose of appropriate memory allocation by a technique of oblivious execution, sketched by Figure 1. 1. Partition the input set into buckets by a random polynomial of constant degree. 2. For t := 1 to O(lg lg n) do (a) Allocate M t memory blocks, each of size K t . (b) Let each bucket select a block at random, and try to injectively map its into the block using a random linear function; if the same block was selected by another bucket, or if no injective mapping was found, then the bucket carries on to the next iteration. Figure 1: The template for the hashing algorithm. The crux of the algorithm is a careful a priori selection of the parameters M t and K t . For each iteration t, M t and K t depend on the expected number of active buckets and the expected distribution of bucket sizes at iteration t in a way that makes the desired progress possible (or rather, likely). The execution is oblivious in the following sense: All buckets are treated equally, regardless of their sizes. The algorithm does not make any explicit attempt to estimate the sizes of individual buckets and to allocate memory to buckets based on their sizes, as is the case in the previous implementations of the 2-level scheme. Nor does it attempt to estimate the number of active buckets or the distribution of their sizes. The selection of the parameters M t and K t in iteration t is made according to a priori estimates of the above random variables. These estimates are based on properties of the level-1 hash function as well as on inductive assumptions about the behavior of previous iterations. Remark The hashing result demonstrates the power of randomness in parallel computation on crcw machines with memory restricted to linear size. Boppana [6] considered the problem of Element Distinctness: given n integers, decide whether or not they are all distinct. He showed that solving Element Distinctness on an n-processor Priority machine with bounded memory requires \Omega\Gammaeq n= lg lg n) time. "Bounded memory" means that the memory size is an arbitrary function of n but not of the range of the input values. It is easy to see that if the memory size is bounded by \Omega\Gamma then Element Distinctness can be solved in O(1) expected time by using hash functions (Fact 2.2). This, however, does not hold for linear size memory. Our parallel hashing algorithm implies that when incorporating randomness, Element Distinctness can be solved in expected O(lg lg n) time using n processors on Collision + (which is weaker than the Priority model) with linear memory size. 1.3 Applications The perfect hash table data structure is a useful tool for parallel algorithms. Matias and Vishkin [39] proposed using a parallel hashing scheme for space reduction in algorithms in which a large amount of space is required for communication between processors. Such algorithms become space efficient and preserve the number of operations. The penalties are in introducing randomization and in having some increase in time. Using our hashing scheme, the time increase may be substantially smaller. There are algorithms for which, by using the scheme of [39], the resulting time increase is O(lg n). By using the new scheme, the time increase is only O (lg lg n lg n). This is the case in the construction of suffix trees for strings [2, 17] and in the naming assignment procedure for substrings over large alphabets [17]. For other algorithms, the time increase in [39] was O(lg lg n) or O (lg lg n) 2 , while our algorithm leaves the expected time unchanged. Such is the case in integer sorting over a polynomial range [33] and over a super-polynomial range [5, 39]. More applications are discussed in the conclusion section. 1.4 Outline The rest of the paper is organized as follows. Preliminary technicalities used in our algorithm and its analysis are given in Section 2. The algorithm template is presented in greater detail in Section 3. Two different implementations, based on different selections of M t and K t , are given in the subsequent sections. Section 4 presents an implementation that does not fully satisfy the statements of Theorem 1 but has a relatively simple analysis. An improved implementation of the main algorithm, with more involved analysis, is presented in Section 5. In Section 6 we show how to reduce the number of random bits. Section 7 explains how the algorithm can be implemented with an optimal number of operations. The model of computation is discussed in Section 8, where we also give a modified algorithm for a weaker model. Section 9 briefly discusses the extension of the hashing problem, in which the input may consist of a multi-set. Finally, conclusions are given in Section 10. Preliminaries The following inequalities are standard (see, e.g. [1]): Markov's inequality Let ! be a random variable assuming non-negative values only. Then Chebyshev's inequality Let ! be a random variable. Then, for T ? 0, Chernoff's inequality Let ! be a binomial variable. Then, for T ? 0, Terminology for probabilities We say that an event occurs with n-dominant probability if it occurs with probability \Gamma\Omega\Gamma21 . Our usage of this notation is essentially as follows. If a poly-logarithmic number of events are such that each one of them occurs with n-dominant probability, then their conjunction occurs with n-dominant probability as well. We will therefore usually be satisfied by demonstrating that each algorithmic step succeeds with n-dominant probability. Fact 2.1 Let independent binary random variables, and let 2. -dominant probability, and 3. Proof. Recall the well known fact that 1-i-n are pairwise independent 1. By Inequality (2) 2. If - 2 - E (!) 2 then by Inequality (2) then by Inequality (1) 3. Follows immediately from the above. Hash functions For the remainder of this section, let S ' U be fixed, splits the set S into buckets; bucket i is the subset fx its size is collides if its bucket is not a singleton. The function is injective, or perfect , if no element collides. Let 0-i!m r A function is injective if and only if is the number of collisions of pairs of keys. More generally, B r is the number of r-tuples of keys that collide under h. Polynomial hash functions Let prime. The class of degree-d polynomial hash functions, d - 1, mapping U into [0; mod m; for some c In the rest of this section we consider the probability space in which h is selected uniformly at random from H d The following fact and corollary were shown by Fredman, Koml'os, and Szemer'edi [16], and before by Carter and Wegman [8]. (The original proof was only for the case however the generalization for d ? 1 is straightforward.) Fact 2.2 E Corollary 2.3 The hash function h is injective on S with probability at least 1 Proof. The function h is injective if and only if Fact 2.2 and Markov's inequality, the probability that h is not injective is Prob (B The following was shown in [11]. Fact 2.4 If d - 3 then For r - 0, let A r be the rth moment of the distribution of s i , 0-i!m s r It is easy to see that A Further, it can be shown that if were completely random function, then A r is linear in n with high probability for all fixed r - 2. For polynomial hash functions, Dietzfelbinger et al. [12] proved the following fact: Fact 2.5 Let r - 0, and m - n. If d - r then there exists a constant oe r ? 0, depending only on r, such that Tighter estimates on the distribution of A r were given in [11]: (For completeness, the proofs are attached in Appendix B.) Fact 2.6 Let r - 2. If d - r then 1-j-r r where \Phi r \Psi is the Stirling number of the second kind. 1 Fact 2.7 Let ffl ? 0 be constant. If d - probability. the Stirling number of the second kind , is the number of ways of partitioning a set of k distinct elements into j nonempty subsets (e.g., [31, Chapter 6]). 3 A Framework for Hashing by Oblivious Execution 3.1 An algorithm template The input to the algorithm is a set S of n keys, given in an array. The hashing algorithm works in two stages, which correspond to the two level hashing scheme of Fredman, Koml'os, and Szemer'edi [16]. In the first stage a level-1 hash function f is chosen. This function is selected at random from the class H d , where d is a sufficiently large constant to be selected in the analysis, and The hash function f partitions the input set into m buckets ; bucket i, is the (i). The first stage is easily implemented in constant time. The main effort is in the implementation of the second stage, which is described next. The second level of the hash table is built in the second stage of the algorithm. For each bucket a private memory region, called a block , is assigned. The address of the memory block allocated to bucket i is recorded in cell i of a designated array ptr of size m. Also, for each bucket, a level-2 function is constructed; this function injectively maps the bucket into its block. The descriptions of the level-2 functions are written in ptr. Let us call a bucket active if an appropriate level-2 function has not yet been found, and inactive otherwise. At the beginning of the stage all buckets are active, and the algorithm terminates when all buckets have become inactive. The second stage consists of O(lg lg n) iterations, each executing in constant time. The iterative process rapidly reduces the number of active buckets and the number of active keys. At each iteration t, a new memory segment is used. This segment is partitioned into M t blocks of size K t each, where M t and K t will be set in the analysis. Each bucket and each key is associated with one processor. The operation of each active bucket in each iteration is given in Figure 2. Allocation: The bucket selects at random one of the M t memory blocks. If the same block was selected by another bucket, then the bucket remains active and does not participate in the next step. Hashing : The bucket selects at random two functions from H 1 , and then tries to hash itself into the block separately by each of these functions. If either one of the functions is injective, then its description and the memory address of the block are written in the appropriate cell of array ptr and the bucket becomes inactive. Otherwise, the bucket remains active and carries on to the next iteration. Figure 2: The two steps of an iteration, based on oblivious execution. In a few of the last iterations, it may become necessary for an iteration to repeat its body more than once, but no more than a constant number of times. The precise conditions and the number of repetitions are given in Section 5. The hash table constructed by the algorithm supports lookup queries in constant time. Given a key x, a search for it begins by reading the cell ptr[f(x)]. The contents of this cell defines the level-2 function to be used for x as well as the address of the memory block in which x is stored. The actual offset in the block in which x is stored is given by the injective level-2 hash function found in the Hashing step above. 3.2 Implementations The algorithm template described above constitutes a framework for building parallel hashing algorithms. The execution of these algorithms is oblivious in the sense that the iterative process of finding level-2 hash functions does not require information about the number or size of active buckets. Successful termination and performance are dependent on the a priori setting of the parameters d, M t and K t . The effectiveness of the allocation step relies on having sufficiently many memory blocks; the effectiveness of the hashing step relies on having sufficiently large memory blocks. The requirement of keeping the total memory linear imposes a tradeoff between the two parameters. The challenge is in finding a balance between M t and K t , so as to achieve a desired rate of decay in the number of active buckets. The number of active keys can be deduced from the number of active buckets based on the characteristics of the level-1 hash function, as determined by d. We will show two different implementations of the algorithm template, each leading to an analysis of a different nature. The first implementation is given in Section 4. There, the parameters are selected in such a way that in each iteration, the number of active buckets is expected to decrease by a constant factor. Although each iteration may fail with constant probability, there is a geometrically decreasing series which bounds from above the number of active buckets in each iteration. After O(lg lg n) iterations, the expected number of active keys and active buckets becomes n=(lg n) \Omega\Gamma1/ . The remaining keys are hashed in additional constant time using a different approach, after employing an O(lg lg n) time procedure. From a technical point of view, the analysis of this implementation imposes relatively modest requirements on the level-1 hash function, since it only uses first-moment analysis (i.e., Markov's inequality). Moreover, it only requires a simpler version of the hashing step, in which only one hash function from H 1 is being used. The expected running time is O(lg lg n), but this running time is guaranteed only with (arbitrary small) constant probability. The second implementation is given in Section 5. This implementation is characterized by a doubly-exponential rate of decrease 2 in the number of active buckets and keys. After O(lg lg n) sequence v0 ; decreases in an exponential rate if for all t, v t - v0=(1 the sequence decreases in a doubly-exponential rate if for all t, v t - v0=2 (1+ffl) t for some ffl ? 0. iterations all keys are hashed without any further processing. This implementation is superior in several other respects: its time performance is with high probability, each key is only handled by its original processor, and it forms a basis for further improvements in reducing the number of random bits. From a technical point of view, the analysis of this implementation is more subtle and imposes more demanding requirements on the level-1 hash function, since it uses second-moment analysis (i.e., Chebyshev's inequality). Achieving a doubly-exponential rate of decrease required a more careful selection of parameters, and was done using a "symbolic spreadsheet" approach. Together, these implementations demonstrate two different paradigms for fast parallel randomized algorithms, each involving a different flavor of analysis. One only requires an exponential rate of decrease in problem size, and then relies on reallocation of processors to items. (Subsequent works that use this paradigm and its extensions are mentioned in Section 10.) This paradigm is relatively easy to understand and not too difficult to analyze, using a framework of probabilistic induction and analysis by expectations. The analysis shows that each iteration succeeds with constant probability, and that this implies an overall constant success probability. In contrast, the second implementation shows that each iteration succeeds with n-dominant probability, and that this implies an overall n-dominant success probability. The analysis is significantly more subtle, and relies on more powerful techniques of second moment analysis. The second paradigm consists of a doubly-exponential rate of decrease in the problem size, and hence does not require any wrap-up step. 4 Obtaining Exponential Decrease This section presents our first implementation of the algorithm template. Using a rather elementary analysis of expectations, we show that at each iteration the problem size decreases by a constant factor with (only) constant probability. The general framework described in Section 4.1 shows that this implies that the problem size decreases at an overall exponential rate. After O(lg lg n) iterations, the number of keys is reduced to n=(lg n)\Omega\Gamma1/ . A simple load balancing algorithm now allocates (lg n)\Omega\Gamma21 processors to each remaining key. Using the excessive number of processors, each key is finally hashed in constant time. 4.1 Designing by Expectation Consider an iterative randomized algorithm, in which after each iteration some measure of the problem decreases by a random amount. In a companion paper [22] we showed that at each iteration one can actually assume that in previous iterations the algorithm was not too far from its expected behavior. The paradigm suggested is: Design an iteration to be "successful" with a constant probability under the assumption that at least a constant fraction of the previous iterations were "successful". It is justified by the following lemma. Lemma 4.1 (probabilistic induction [22]) Consider an iterative randomized process in which, for all t - 0, the following holds: iteration t with probability at least 1=2, provided that among the first t iterations at least t=4 were successful. Then, with probability\Omega\Gammail , for every t ? 0 the number of successful iterations among the first t iterations is at least t=4. 4.2 Parameters setting and analysis Let the level-1 function be taken from H 10 Further, set Let is as in Fact 2.5. To simplify the analysis, we allow the parameters K t and M t to assume non-integral values. In actual implementation, they must be rounded up to the nearest integer. This does not increase memory requirements by more than a constant factor; all other performance measures can only be improved. Memory usage The memory space used is Lemma 4.2 Let v t be the number of active buckets at the beginning of iteration t. Then, Proof. We assume that the level-1 function f satisfies By Fact 2.5, (11) holds with probability at least 1=2. The proof is by continued by using Lemma 4.1. Iteration t is successful if v t+1 - v t =2. Thus, the number of active buckets after j successful iterations is at most m2 \Gammaj . The probabilistic inductive hypothesis is that among the first t iterations at least t=4 were successful, that is The probabilistic inductive step is to show that In each iteration the parameters K t and M t were chosen so as to achieve constant deactivation probability for buckets of size at most We distinguish between the following three types of events, "failures", which may cause a bucket to remain active at the end of an iteration. (i) Allocation Failure. The bucket may select a memory block which is also selected by other buckets. be the probability that a fixed bucket does not successfully reserve a block in the allocation step. Since there are at most v t buckets, each selecting at random one of M t memory blocks, ae 1 (t) - v t =M t . By (12) and (10) (ii) Size Failure. The bucket may be too large for the current memory block size. As a result, the probability for it to find a level-2 hash function is not high enough. t be the number of buckets at the beginning of iteration t that are larger than fi t . By (11), Therefore, by (13), Without loss of generality, we assume that if v t+1 - v t =2 then v buckets that are needed become inactive, then some of them are still considered as active). Thus, for the purpose of analysis, We have then (iii) Hash Failure. A bucket may fail to find an injective level-2 hash function even though it is sufficiently small and it has uniquely selected a block. Let ae 3 (t) be the probability that a bucket of size at most fi t is not successfully mapped into a block of size K t in the hashing step. By Corollary 2.3 and (13) A bucket of size at most fi t that successfully reserves a block of size K t , and that is successfully mapped into it, becomes inactive. The expected number of active buckets at the beginning of iteration therefore be bounded by By Markov's inequality proving the inductive step. The lemma follows. Lemma 4.3 Let n t be the number of active keys at the beginning of iteration t. Then for some constants c; ff ? 0. Proof. It follows from (11), by using a simple convexity argument, that n t is maximal when all active buckets at the beginning of iteration t are of the same size q t . In this case, by (11), and Therefore, by Lemma 4.2, the lemma follows. By Lemma 4.3 and Lemma 4.2 we have an exponential decrease in the number of active keys and in the number of active buckets with probability\Omega\Gamma329 The number of active keys becomes n=(lg n) c , for any constant c ? 0, after O(lg lg n) iterations with 4.3 A final stage After the execution of the second stage with the parameter setting as described above, the number of available resources (memory cells and processors) is a factor of (lg n)\Omega\Gamma1/ larger than the number of active keys. This resource redundancy makes it possible to hash the remaining active keys in constant time, as described in the remainder of this section. All keys that were not hashed in the iterative process will be hashed into an auxiliary hash table of size O(n). Consequently, the implementation of a lookup query will consist of searching the key in both hash tables. The auxiliary hash table is built using the the 2-level hashing scheme. A level-1 function maps the set of active keys into an array of size n. This function is selected at random from a class of hash functions presented by Dietzfelbinger and Meyer auf der Heide [14, Definition 4.1]. It has the property that with n-dominant probability each bucket is of size smaller than lg n [14, Theorem 4.6(b')]. For the remainder of this section we assume that this event indeed occurs. (Alternatively, we can use the n ffl -universal class of hash functions presented by Siegel [43].) Each active key is allocated 2 lg n processors, and each active bucket is allocated 4(lg n) 3 mem- ory. The allocation is done by mapping the active keys injectively into an array of size O(n= lg n), and by mapping the indices of buckets injectively into an array of size O(n=(lg n) 3 ). These mappings can be done in O(lg lg n) time with n-dominant probability, by using the simple renaming algorithm from [20]. The remaining steps take constant time. We independently select 2 lg n linear hash functions and store them in a designated array. These hash functions will be used by all buckets. The memory allocated to each bucket is partitioned into 2 lg n memory blocks, each of size 2 lg 2 n. Each bucket is mapped in parallel into its 2 lg n blocks by the 2 lg n selected linear hash functions, and each mapping is tested for injectiveness. This is carried out by the 2 lg n processors allocated to each key. For each bucket, one of the injective mappings is selected as a level-2 function. The selection is made by using the simple 'leftmost 1' algorithm of [15]. If for any of the buckets all the mappings are not injective then the construction of the auxiliary hash table fails. Lemma 4.4 Assume that the number of keys that remain active after the iterative process is at most n=(lg n) 3 . Then, the construction of the auxiliary hash table succeeds with n-dominant probability. Proof. Recall that each bucket is of size at most lg n; A mapping of a bucket into its memory block of size 2(lg n) 2 is injective with probability at least 1=2 by Corollary 2.3. The probability that a bucket has no injective mapping is therefore at most 1=n 2 . With probability at least 1 \Gamma 1=n, every bucket has at least one injective mapping. It is easy to identify failure. If the algorithm fails to terminate within a designated time, it can be restarted. The hash table will be therefore always constructed. Since the overall failure probability is constant, the expected running time is O(lg lg n). 5 Obtaining Doubly-Exponential Decrease The implementation of the algorithm template that was presented in the previous section maintains an exponential decrease in the number of active buckets throughout the iterations. This section presents the implementation in which the number of active buckets decreases at a doubly- exponential rate. Intuitively, the stochastic process behind the algorithm template has a potential for achieving doubly-exponential rate: If a memory block is sufficiently large in comparison to the bucket size then the probability of the bucket to remain active is inversely proportional to the size of the memory block (Corollary 2.3). Consider an idealized situation in which this is the case. If at iteration t there are m t active buckets, each allocated a memory block of size K t , then at iteration t will be m t =K t active buckets, and each of those could be allocated a memory block of size K 2 t ; at iteration there will be m t =K 3 active buckets, each to be allocated a memory block of size K 4 and so on. In a less idealized setting, some buckets do not deactivate because they are too large for the current value of K t . The number of such buckets can be bounded above by using properties of the level-1 hash function. It must be guaranteed that the fraction of "large buckets" also decreases at a doubly-exponential rate. The illustrative crude calculation given above assumes that memory can be evenly distributed between the active buckets. To make the doubly-exponential rate possible, the failure probability of the allocation step, and hence the ratio m t =M t , must also decrease at a doubly-exponential rate. Establishing a bound on the number of "large blocks" and showing that a large fraction of the buckets are allocated memory blocks were also of concern in the previous section. There, however, it was enough to show constant bounds on the probabilities of allocation failure, size failure and hash failure. The parameter setting which establishes the balance required for the doubly-exponential rate is now presented. Following that is the analysis of the algorithm performance. The section concludes with a description of how the parameters were selected. 5.1 Parameters setting Let the level-1 function be taken from H Further, set Let where 5.2 Memory usage Proposition 5.1 The total memory used by the algorithm is O(n). Proof. By (17), the memory used in the first stage is O(n). The memory used in an iteration t of the second stage is The total memory used by the second stage is therefore at mostX 5.3 Framework for time performance analysis be defined by The run-time analysis of the second stage is carried out by showing: Lemma 5.2 With n-dominant probability, the number of active buckets in the beginning of iteration t is at most m t . The lemma is proved by induction on t, for t - lg lg n= lg -. The induction base follows from and the fact that there are at most n active buckets. In the subsequent subsections, we prove the inductive step by deriving estimates on the number of failing buckets in iteration t under the assumption that at the beginning of the iteration there are at most m t active buckets. Specifically, we show by induction on t that, with n-dominant probability, the number of active buckets at the end of iteration t is at most The bucket may fail to find an injective level-2 hash function. In estimating the number of buckets that fail to find an injective level-2 function during an iteration we assume that the bucket uniquely selected a memory block and that the bucket size is not too large relatively to the current block size. Accordingly, as in Section 4.2, we distinguish between the following three types of events, "failures", which may cause a bucket to remain active at the end of an iteration. (i) Allocation Failure. The bucket may select a memory block which is also selected by other buckets. (ii) Size Failure. The bucket may be too large for the current memory block size. As a result, the probability for it to find a level-2 hash function is not high enough. (iii) Hash Failure. A bucket may fail to find a level-2 hash function even though it is sufficiently small and it has uniquely selected a block. We will provide estimates for the number of buckets that remain active due to either of the above reasons: in Lemma 5.5 for case (i), in Lemma 5.6 for case (ii), and in Lemma 5.7 and Lemma 5.8 for case (iii). The estimates are all shown to hold with n-dominant probability. The induction step follows from adding all these estimates. To wrap up, let We can therefore infer: Proposition 5.3 With n-dominant probability, the number of iterations required to deactivate all buckets is at most lg lg n= lg -. 5.4 Failures in Uniquely Selecting a Block Lemma 5.4 Let ffl be fixed, suppose that either m t ? M 1=2+ffl t . Let ! be the random variable representing the number of buckets that fail to uniquely select a block. Proof. A bucket has a probability of at most m t =M t to have other buckets select the memory block it selected. Therefore, Further, ! is stochastically smaller than a binomially distributed random variable $ obtained by performing m t independent trials, each with probability m t =M t of success. That is to say, t then by (3) \Gamma\Omega\Gamma26 Otherwise, t and we are in the situation where E (!) - 1. Since ! is integer valued and 2m 2 by (1) by (25) The setting not covered by the above lemma is M 1=2\Gammaffl t . This only occurs in a constant number of iterations throughout the algorithm and requires the following special treatment. The body of these iterations is repeated, thus providing a second allocation attempt of buckets that failed to uniquely select a memory block in the first trial. be the random variables representing the number of buckets that fail to uniquely select a block in the first and second attempts respectively. j by (1) by (25) \Gamma\Omega\Gamma27 Therefore, with M t -dominant probability the second attempt falls within the conditions of Equation Lemma 5.5 Let t - lg lg n= lg -. The number of buckets that fail to uniquely select a block is, with n-dominant probability, at most m t+1 =4. Proof. By Lemma 5.4, the number of buckets that fail to uniquely select a memory block is, with -dominant probability, at most by (19),(23) by (20) by (20) by (24) The above holds also with n-dominant probability since by (19) by (20) 5.5 Failures in Hashing In considering buckets which uniquely selected a block which fail to find an injective level-2 function we draw special attention to buckets of size at most Lemma 5.6 The number of buckets larger than fi t is, with n-dominant probability, at most m t+1 =4. Proof. Let incorporating the appropriate values for the Stirling numbers of the second kind into Fact 2.6, we get by (17) Therefore, by Fact 2.7, with n-dominant probability From the above and (6) it follows that the number of buckets bigger than fi t is, with n-dominant probability, at most by (31),(16) by (18) by (20) by (20) by (20) by (24) The analysis of hashing failures of buckets that are small enough is further split into two cases. Lemma 5.7 Suppose that m t =2K t - n. Then the number of buckets of size at most fi t that fail in the hashing step of the iteration is, with n-dominant probability, at most m t+1 =4. Proof. Without loss of generality, we may assume that there are exactly m t active buckets of size at most fi t that participate in Step 2. When such a bucket is mapped into a memory block of size K t , the probability of the mapping being non-injective is, by Corollary 2.3, at most fi 2 The probability that the bucket fails in both hashing attempts is therefore at most 1=2K t . Let ~ t be the total number of such failing buckets. Then, . By Fact 2.1, with ~ by (18),(23) by (20) by (20) by (20),(24) Note that since m t =2K t - n, the above holds with n-dominant probability and we are done. Lemma 5.8 Suppose that m t n. Then, by repeating the hashing step of the iteration a constant number of times, we get ~ Proof. We have and thus, Therefore, by (18) by (36) for some constant ffi ? 0. Recall from the proof of Lemma 5.7 that a bucket fails in the hashing step with probability at most 1=2K t . By (37), if the iteration body is repeated d2=ffie the failure probability of each bucket becomes at most (2K t The lemma follows by Markov inequality. 6 Reducing the Number of Random Bits In this section we show how to reduce the number of random bits used by the hashing algorithm. The algorithm as described in the previous section consumes \Theta(n lg u) random bits, where the first iteration already uses \Theta(n lg u) random bits; for each subsequent iteration, the number of random words from U which are used is by at most a constant factor larger than the memory used in that iteration, resulting in a total of \Theta(n lg u) random bits. The sequential hashing algorithm of Fredman, Koml'os, and Szemer'edi [16] can be implemented with only O(lg lg U bits [11]. We show how the parallel hashing algorithm can be implemented with O(lg lg U bits. We first show how the algorithm can be modified so as to reduce the number of random bits to O(lg u lg lg n). The first stage requires O(1) random elements from U for the construction of the level-1 function, and remains unchanged. An iteration t of the second stage required O(m t ) random elements from U ; it is modified as follows. Allocation step If each bucket independently selects a random memory block then O(m t lg M t ) random bits are consumed. This can be reduced to O(lg m) by making use of polynomial hash functions Lemma 6.1 Using 6 lg m random bits, a set R ' t can be mapped in constant time into an array of size 3M t such that the number of colliding elements is at most 2m 2 Proof. Let and be selected at random. Then, the image of a bucket i is defined by Algorithmically, h 1 is first applied to all elements and then h 2 is applied to the elements which collided under h 1 . The colliding elements of g t are those which collided both under h 1 and under h 2 . R 0 be the set of elements that collide under h 1 . Clearly, jR 1=6. Consider the following three cases: By Corollary 2.3, Prob (R 0 6= 2. It follows from Fact 2.4 that probability. As t =2 we have that jR 3. M 1=2\Gammaffl By Fact 2.2, t =2 and by Markov's inequality, Therefore, with M t -dominant probability, jR t , in which case, by Corollary 2.3, is not injective over R 0 Invoking the above procedure for block allocation does not increase the total memory consumption of the algorithm by more than a constant factor. Hashing step The implementation of the hashing part of the iteration body using independent hash functions for each of the active buckets consumes O(m t lg u) random bits. This can be reduced to O(lg u) by using hash functions which are only pairwise independent . This technique and its application in the context of hash functions are essentially due to [10, 11]. The modification to the step is as follows. In each hashing attempt executed during the step, four global parameters a are selected at random by the algorithm. The hash function attempted by a bucket i is where All hashing attempts of the same bucket are fully independent. Thus, the proof of Lemma 5.8 is unaffected by this modification. Recall that Fact 2.1 assumes only pairwise independence. Since are pairwise independent, the proof of Lemma 5.7 remains valid as well. The above leads to a reduction in the number of random bits used by the algorithm to O(lg u lg lg n). The number of random bits can be further reduced as follows: Employ a pre-processing hashing step in which the input set S is injectively mapped into the range [0; This is done by applying a hash function - selected from an appropriate class, to map the universe U into this range. Then the algorithm described above is used to build a hash table for the set -(S). A lookup of a key x is done by searching for -(x) in this hash table. The simple class of hash functions H 3 m is appropriate for this universe reduction application. It was shown in [11] that the class H 3 m has the following properties: 1. A selection of a random function - from the class requires O(lg lg u bits. 2. A selection can be made in constant time by a single processor. 3. The function - is injective over S with n-dominant probability. 4. Computing -(x) for any x 2 U can be done in constant time. This pre-processing is tantamount to a reduction in the size of the universe, after which application of the algorithm requires only O(lg n lg lg n) bits. The total number of random bits used is therefore O(lg lg u 7 Obtaining Optimal Speedup The description of the algorithm in Section 3 assumed that the number of processors is n; thus the time-processor product is O(n lg lg n). Our objective in this section is a work-optimal implementation where this product is O(n), and p, the number of processors, is maximized. array and the bucket array are divided into p sectors , one per processor. A parallel step of the algorithm is executed by having each processor traverse its sector and execute the tasks included in it. A key is active if its bucket is active. Let n t be the number of active keys in the beginning of iteration t. Assume that the implemented algorithm has reached the point where Further assume that these active elements are gathered in an array of size O(n= lg lg n). Then, applying the non-optimal algorithm of Section 3 with p - n= lg lg n, and each processor being responsible for n=p lg lg n problem instances, gives a running time of O lg lg n O (n=p) which is work-optimal. We first show that the problem size is reduced sufficiently for the application of the non-optimal algorithm after O(lg lg lg lg n) iterations. Lemma 7.1 There exists t n) such that n t n) with n-dominant probability Proof. The number of active buckets decreases at a doubly-exponential rate as can be seen from Lemma 5.2. To see that the number of keys decreases at a doubly-exponential rate as well, we show that with n-dominant probability Inequality (32), A r - 6n, clearly holds when the summation is over active buckets only. By a convexity argument, the total number of keys in active buckets is maximized when all active buckets are of equal size. The number of active buckets is bounded from above by m t . Therefore, Inequality (40) is obtained from (41) by replacing in m t by its definition in (23) and then substituting numerical values for the parameters using (16) and (20). The lemma follows by choosing an appropriate value for t 0 with respect to (23) and (40). It remains to exhibit a work-efficient implementation of the first t 0 steps of the algorithm. This implementation outputs the active elements gathered in an array of size O(n= lg lg n). The rest of this section is dedicated to the description of this implementation. As the algorithm progresses, the number of active keys and the number of active buckets de- crease. However, the decrease in the number of active elements in different sectors is not necessarily identical. The time of implementing one parallel step is proportional to the number of active elements in the largest sector. It is therefore crucial to occasionally balance the number of active elements among different sectors in order to obtain work efficiency. Let the load of a sector be the number of active elements (tasks) in it. A load balancing algorithm takes as input a set of tasks arbitrarily distributed among p sectors; using p processors it redistributes this set so that the load of each sector is greater than the average load by at most a constant factor. Suppose that we have a load balancing algorithm whose running time, using p processors, is T lb (p) with n-dominant probability. If load balancing is applied after step t then the size of each sector is O(n t =p). We describe a simple work-optimal implementation in which load balancing is applied after each of the first t 0 parallel steps. A parallel step t executes in time which is in the order of The total time of this implementation is in the order of decreases at least at an exponential rate, the total time is in the order of which is O(n=p) for Using the load balancing algorithm of [20] which runs in T lb time, we conclude that with n-dominant probability the running time on a p-processor machine is The load balancing algorithm applied consumes O(p lg lg p) random bits. All these bits are used in a random mapping step which is very similar to the allocation step of the hashing algorithm. Thus, by a similar approach as the mapping procedure in Lemma 6.1 it may be established that the number of random bits in the load balancing algorithm can be reduced to O(lg p lg lg p). We finally remark that using load balancing in a more efficient, yet as simple way, as describe in [23], yields a faster work-efficient implementation. The technique is based on carefully choosing the appropriate times for invoking the load balancing procedure; it applies to any algorithm in which the problem size has an exponential rate of decrease, and it hence applies to the implementation of Section 4 as well. In such an implementation the load balancing algorithm is only used O(lg n) times, resulting in a parallel hashing algorithm that takes O(n=p+lg lg n lg n) time with n-dominant probability. 8 Model of Computation In this section we give a closer attention to the details of the implementation on a pram, and study the type of concurrent memory access required by our algorithm. We first present an implementation on Collision, and its extension to the weaker Tolerant model. We proceed by presenting an implementation on the even weaker Robust model. The hash-table constructed in this implementation only supports searches in O(lg lg n) time. Finally, we examine the concurrent read capability needed by the implementations. 8.1 Implementation on Collision and on Tolerant We describe an implementation on Collision. This implementation is also valid for Tolerant, since each step of Collision can be simulated in constant time on Tolerant provided that, as it is the case here, only linear memory is used [32]. Initialization The selection of the level-1 hash function is done by a single processor. Since the level-1 function is a polynomial of a constant degree, its selection can be done by a single processor and be read by all processors in constant time, using a singe memory cell of dmax flg lg u; lg nge bits. No concurrent-write operation is required for the implementation of this stage. Bucket representatives The algorithm template assumes that each bucket can act as a single entity for some operations, e.g., selecting a random block and selecting a random hash function. Since usually several keys belong to the same bucket, it is necessary to coordinate the actions of the processors allocated to these keys. A simple way of doing so is based on the fact that there are only linearly many buckets and that a bucket is uniquely indexed by the value of f , the level-1 hash function, on its members. A processor whose index is determined by the bucket index acts as the bucket representative and performs the actions prescribed by the algorithm to the bucket. Allocation and Hashing steps A processor representing an active bucket selects a memory block and a level-2 hash function, and records these selections in a designated cell. All processors with keys in that bucket read then that cell and use the selected block in the hashing step. Each participating processor (whose key belongs in an active bucket) writes its key in the cell determined by its level- examines the cell contents to see if the write operation was successful. A processor for which the write failed will then attempt to write its key to position i of array ptr, where i is the number of the bucket this processor belongs to. Processors belonging to bucket i can then learn if the level-2 function selected for their bucket is injective by reading the content of ptr[i]. A change in value or a collision symbol indicate non-injectiveness. To complete the process, the array ptr is restored for the next hashing attempt. This restoration can be done in constant time since this array is of linear size. In summary we have Proposition 8.1 The algorithms of Theorem 1 can be implemented on Tolerant. 8.2 Implementation on Robust We now describe an implementation that, at the expense of slowing down the lookup operation, makes no assumption about the result of a concurrent-write into a cell. Specifically, we present an implementation on the Robust model, for which a lookup query may take O(lg lg n) time in the worst case, but O(1) expected time for keys in the table. The difficulty with the Robust model is in letting all processors in a bucket know whether the level-2 hash function of their bucket is injective or not. The main idea in the modified implementation is in allowing iterations to proceed without determining whether level-2 hash functions are injective or not; whenever a key is written into a memory cell in the hashing step it is deactivated, and its bucket size decreases. The modified algorithm performs at least as well as the implementation in which a bucket is deactivated only if all of its keys are mapped injectively. The total memory used by the modified algorithm and the size of the representation of the hash table do not change. Allocation step We first note that the algorithm can be carried out without using bucket representatives at all. Allocation of memory blocks is done using hash functions, as in Lemma 6.1; each processor can individually compute the index of its memory block by evaluating the function g t . This function is selected by a designated processor and its representation (6 lg m bits) is read in constant time by all processors. We further modify the algorithm, so that the hashing step is carried out by all active buckets. That is, even buckets that collided in the allocation step will participate in the hashing step. This modification can only serve to improve the performance of the algorithm, since even while sharing a block with another bucket the probability that a bucket finds an injective function into that block is not zero. This modification eliminates the concurrent memory access needed for detecting failures in the allocation step. Hashing step The selection of a level-2 hash function is done as in the hashing step described in Section 6. As can be seen from (39), only four global parameters should be selected and made available to all processors; this can be done in constant time. It remains to eliminate the concurrent memory access required for determining if the level-2 function of any single bucket was injective. Whenever a key is successfully hashed by this function, it is deactivated even if other keys in the same bucket were not successfully hashed. Thus, keys of the same bucket may be stored in the hash table using different level-2 hash functions. The two steps of an iteration in the hashing algorithm are summarized in Figure 3. Let x be an active key in a bucket f(x). The processor assigned to x executes the following steps. Allocation: Compute (i), the index of the memory block selected to the bucket of x, where g t is defined by (38). Hashing : Determine h i , the level-2 hash function selected by the bucket of x, where h i is defined by (39). Write x into cell h i (x) in memory block g t (i) and read the contents of that cell; if x was written then the key x becomes inactive. Figure 3: Implementation of iteration t in the hashing algorithm on Robust Lookup algorithm The search for a key x is done as follows. Let read position h i (x) in the memory block g t (i) in the appropriate array. (All random bits that were used in the hash table construction algorithm are assumed to be recorded and available.) The search is terminated when either x is found, or else when t exceeds the number of iterations in the construction algorithm. The lookup algorithm requires O(lg lg n) iterations in the worst case. However, for any key x 2 S the expected lookup time (over all the random selections made by the hashing algorithm) is O(1). An alternative simplified implementation Curiously, the sequence of modifications to the algorithm described in this section has lead to a 1-level hashing scheme, i.e., to the elimination of indirect addressing. To see this, we observe that at iteration t an active key x is written into a memory cell g t (x), where the function g t (x) is dependent only on n and on the random selections made by the algorithm, but not on the input. An even simpler implementation of a 1-level hashing algorithm is delineated next. At each iteration t, a new array T t of size 3M t is used, where M t is as defined in (19). In addition, a function g t as defined in (38) is selected at random. A processor representing an active key x in the iteration tries to write x into T t [g t (x)], and then reads this cell. If x is successfully written in T t [g t (x)] then x is deactivated. Otherwise, x remains active and the processor representing it carries on to the next iteration. To see that the algorithm terminates in O(lg lg n) iterations, we observe that the operation on keys in each iteration is the same as the operation on buckets in the allocation step of Section 6. Therefore, the analysis of Section 6 can be reused, substituting keys for buckets (and ignoring failures in the hashing step of the 2-level algorithm). The hash table consists of the collection of the arrays T and, as can be easily verified, is of linear size. A lookup query for a given key x is executed in O(lg lg n) time by reading T t [g t (x)] for 8.3 Minimizing concurrent read requirements The algorithms for construction of the hash table on Tolerant and Robust can be modified to use concurrent-read from a single cell only. By allowing a pre-processing stage of O(lg n) time, concurrent read can be eliminated, implying that the ercw model is sufficient. With these modifications, parallel lookups still require concurrent read, and their execution time increases to O(lg lg n) in the worst case. Nevertheless, the expected time for lookup of any single key x 2 S is O(1). The details are described next. 8.3.1 Concurrent read in the Tolerant implementation There are two types of concurrent read operations required by the modified algorithm. First, the sequence of O(lg lg n) functions g t (or alternatively, g t in the simplified implementation), must be agreed upon by all processors. Since each of these functions is represented by O(lg u) bits, its selection can be broadcasted at the beginning of the iteration through the concurrent-read cell. The single cell concurrent read requirement for broadcasting can be eliminated by adding an O(lg n)-time pre-processing step for the broadcasting. (This is just a special case of simulating crcw pram by erew pram.) The other kind of concurrent-read operation occurs when processors read a memory cell to verify that their hashing into that cell has succeeded. This operation can be replaced by the following procedure. For each memory cell, there is a processor standing by. Whenever a pair hx; ji is written into a cell, the processor assigned to that cell sends an acknowledgement to processor j by writing into a memory cell j in a designated array. The lookup algorithm requires concurrent-read capabilities. In this sense, the lookup operation is more demanding than the construction of the hash table. A similar phenomenon was observed by Karp, Luby and Meyer auf der Heide [36] in the context of simulating a random access machine on a distributed memory machine. The main challenge in the design of their (parallel-hashing based) simulation algorithm was the execution of the read step. Congestions during the execution of the write step were resolved by attempting to write in several locations and using the first for which the write succeeded. It is more difficult to resolve read congestions since the cells in which values were stored are already determined. Indeed, the read operation constitutes the main run-time bottleneck in their algorithm. 8.3.2 Concurrent Read in the Robust implementation The simplified 1-level hashing algorithm for construction of the hash table on Robust is modified as follows. We eliminate the step in which a processor with key x reads the contents of the cell after trying to write to that cell. Instead, we use the acknowledgement technique described above: A processor j handling an active key x writes hx; ji into the cell T t [g t (x)]. The processor standing by cell T t [g t (x)] into which hx; ji is written, sends an acknowledgement to processor j. Note that this implementation introduces a new type of failures: due to the unpredictability of the concurrent write operation in Robust, an acknowledgement for a successful hash may not be received. Consider for example the following situation: Let j be a processor whose key x did not collide. Let i, i 0 be two processors with colliding keys y, y 0 , i.e., g t two processors concurrently write the pairs hy; ii and hy into the cell T t [g t (y)]. The result of this concurrent write is arbitrary. In particular, it can be the pair hx 0 ; ji, which would cause the processor standing by the cell T t [g t (y)] to garble the acknowledgement sent to processor j. (Recall that an acknowledgement to processor j is implemented by writing into a memory location associated with j.) The number of the new failures described above can be at most half the number of colliding keys. It is easy to verify that the analysis remains valid, since the number of these new failures in no more than the number of "hashing failures" accounted for in Section 5.5, and which do not occur in this implementation. 9 Hashing of Multi-Sets We conclude the technical discussion by briefly considering a variation of the hashing problem in which the input is a multi-set rather than a set. We first note that the analyses of exponential and doubly-exponential rate of decrease in the problem size is not affected by the possibility of multiple occurrences of the same key. This is a result of relying on estimates of the number of active buckets rather than the number of active keys. The number of distinct keys-not the number of keys-determines the probability of a bucket to find an injective function. A predictable decrease in the number of active keys is essential for obtaining an optimal speedup algorithm. Unfortunately, the analysis in Section 7 with regard to the implementation of Section 5 does not hold. To understand the difficulty, consider the case where a substantial fraction of the input consists of copies of the same key. Then, with non-negligible probability this key may belong to a large bucket. The probability that this bucket deactivates in the first few iterations, in which the memory blocks are not sufficiently large, is too small to allow global decrease in the number of keys with high probability. Consequently, the rapid decrease in the number of buckets may not be accompanied by a similar decrease in the number of keys. In contrast, the nature of the analysis in Section 4 makes it susceptible to an easy extension to multiple keys, which leads to an optimal speedup algorithm, albeit with expected performance only. Using the probabilistic induction lemma all that is required is to show that each copy of an active key stands a constant positive probability of deactivation at each iteration. Since the analysis is based on expectations only, there are no concerns regarding correlations between copies of the same key, or dependencies between different iterations. The details are left to the reader. We also note that the model of computation required for a multi-set is Collision + , since it must be possible to distinguish between the case of multiple copies of the same key being written into a memory cell, and the case where distinct keys are written. Also, the extensions of the hashing algorithms which only require concurrent read from a single memory cell can be used for hashing with multi-set input, but then a Collision + model, as opposed to Robust, must be assumed. We finally observe that the hashing problem with a multi-set as input can be reduced into the ordinary hashing problem (in which the input consists of a set), by a procedure known as leaders election. This procedure selects a single representative from among all processors which share a value. By using an O(lg lg n)-time, linear-work leaders election algorithm which runs on Tolerant [24] we have Theorem 2 Given a multi-set of n keys drawn from a universe U , the hashing problem can be solved using O(n) space: (i) in O(lg lg n) time with high probability, using n processors, or (ii) in O(lg lg n lg n) time and O(n) operations with high probability. The algorithms run on Tolerant. Conversely, note that any hashing algorithm, when run on Arbitrary, solves the leaders election problem. In particular, the simple 1-level hashing algorithm for Robust, when implemented on Arbitrary with a multi-set as input, gives a simple leaders election algorithm. Consider now another variant of the multi-set hashing problem in which a data record is associated with each key. The natural semantics of this problem is that multiple copies of the same key can be inserted into the hash table only if their data records are identical. Processors representing copies of a key with conflicting data records should terminate the computation with an error code. The Collision + model makes it easy enough to extend the implementations discussed above to accommodate this variant. A more sophisticated semantics, in which the data records should be consolidated, requires a different treatment, e.g., by applying an integer sorting algorithm on the hashed keys (see [39]). Conclusions We presented a novel technique of hashing by oblivious execution. By using this technique, algorithms for constructing a perfect hash table which are fast, simple, and efficient, were made possible. The running time obtained is best possible in a model in which keys are only handled in their original processors. The number of random bits consumed by the algorithm is \Theta(lg lg u n). An open question is to close the gap between this number and the \Theta(lg lg u bits that are consumed in the sequential hashing algorithm of [11]. The program executed by each processor is extremely simple. Indeed, the only coordination between processors is in computing the and function, when testing for injectiveness. In the implementation on the Robust model, even this coordination is eliminated. The large constants hidden under the "Oh" notation in the analysis may render the described implementations still far from being practical. We believe that the constants can be substantially improved without compromising the simplicity of the algorithm, by a more careful tuning of the parameters and by tightened analysis. This may be an interesting subject of a separate research. The usefulness of the oblivious execution approach presented in this paper is not limited to the hashing problem alone. We have adopted it in [24] for simulations among sub-models of the crcw pram. As in the hashing algorithm, keys are partitioned into subsets. However, this partition is arbitrary and given in the input, and for each subset the maximum key must be computed. Subsequent work The oblivious execution technique for hashing from Section 3 and its implementation from Section 4 were presented in preliminary form in [21]. Subsequently, our oblivious execution technique was used several times to obtain improvements in running time of parallel hashing algorithms: Matias and Vishkin [38] gave an O(lg n lg lg n) expected time algorithm; Gil, Matias, and Vishkin [26] gave a tighter failure probability analysis for the algorithm in [38], yielding O(lg n) time with high similar improvement (from O(lg n lg lg n) expected time to O(lg n) time with high probability), was described independently by Bast and Hagerup [3]. An O(lg n) time hashing algorithm is used as a building block in a parallel dictionary algorithm presented in [26]. parallel dictionary algorithm supports in parallel batches of operations insert , delete, and lookup.) The oblivious execution technique has an important role in the implementation of insertions into the dictionary. The dictionary algorithm runs in O(lg n) time with high probability, improving the O(n ffl dictionary algorithm of Dietzfelbinger and Meyer auf der Heide [13]. The dictionary algorithm can be used to obtain a space efficient implementation of any parallel algorithm, at the cost of a slowdown of at most O(lg n) time with high probability. The above hashing algorithms use the log-star paradigm of [38], relying extensively on processor reallocation, and are not as simple as the algorithm presented in this paper. Moreover, they require a substantially larger number of random bits. Karp, Luby and Meyer auf der Heide [36] presented an efficient simulation of a pram on a distributed memory machine in the doubly-logarithmic time level, improving over previous simulations in the logarithmic time level. The use of a fast parallel hashing algorithm is essential in their result; the algorithm presented here is sufficient to obtain it. Goldberg, Jerrum, Leighton and Rao [28] used techniques from this paper to obtain an O(h lg lg n) randomized algorithm for the h-relation problem on the optical communication parallel computer model. Gibbons, Matias and Ramachandran [18] adapted the algorithm presented here to obtain a low-contention parallel hashing algorithm for the qrqw pram model [19]; this implies an efficient hashing algorithm on Valiant's bsp model, and hence on hypercube-type non-combining networks [44]. Acknowledgments We thank Martin Dietzfelbinger and Faith E. Fich for providing helpful comments. We also wish to thank Uzi Vishkin and Avi Wigderson for early discussions. Part of this research was done during visits of the first author to AT&T Bell Laboratories, and of the second author to the University of British Columbia. We would like to thank these institutions for supporting these visits. Many valuable comments made by two anonymous referees are gratefully acknowledged. --R The Probabilistic Method. Parallel construction of a suffix tree. Fast and reliable parallel hashing. Optimal bounds for decision problems on the CRCW PRAM. Improved deterministic parallel integer sorting. Optimal separations between concurrent-write parallel machines Observations concerning synchronous parallel models of computation. Universal classes of hash functions. New simulations between CRCW PRAMs. On the power of two-point based sampling Polynomial hash functions are reliable. Relations between concurrent-write models of parallel computation Storing a sparse table with O(1) worst case access time. Data structures and algorithms for approximate string matching. Efficient low-contention parallel algo- rithms The QRQW PRAM: Accounting for contention in parallel algorithms. Fast load balancing on a PRAM. Fast hashing on a PRAM-designing by expectation Designing algorithms by expectations. An effective load balancing policy for geometric decaying algorithms. Fast and efficient simulations among CRCW PRAMs. Simple fast parallel hashing. Towards a theory of nearly constant time parallel algorithms. Doubly logarithmic communication algorithms for optical communication parallel computers. A universal interconnection pattern for parallel computers. Optimal parallel approximation algorithms for prefix sums and integer sorting. Concrete Mathematics. Incomparability in parallel computation. Towards optimal parallel bucket sorting. Every robust CRCW PRAM can efficiently simulate a Priority PRAM. Introduction to Parallel Algorithms. Sorting and Searching Converting high probability into nearly-constant time-with applications to parallel hashing On parallel hashing and integer sorting. Data Structures and Algorithms I: Sorting and Searching. Data structures. An O(lg n) parallel connectivity algorithm. On universal classes of fast high performance hash functions General purpose parallel architectures. --TR

data structures;randomization;parallel computation;hashing

290221

Downward Separation Fails Catastrophically for Limited Nondeterminism Classes.

The $\beta$ hierarchy consists of classes $\beta_k={\rm NP}[logkn]\subseteq {\rm NP}$. Unlike collapses in the polynomial hierarchy and the Boolean hierarchy, collapses in the $\beta$ hierarchy do not seem to translate up, nor does closure under complement seem to cause the hierarchy to collapse. For any consistent set of collapses and separations of levels of the hierarchy that respects ${\rm \cdots \subseteq {\rm NP}$, we can construct an oracle relative to which those collapses and separations hold; at the same time we can make distinct levels of the hierarchy closed under computation or not, as we wish. To give two relatively tame examples: for any $k \geq 1$, we construct an oracle relative to which \[ {\rm \cdots \] and another oracle relative to which \[ {\rm We also construct an oracle relative to which

Introduction . Although standard nondeterministic algorithms solve many NP-complete problems with O(n) nondeterministic moves, there are other problems that seem to require very different amounts of nondeterminism. For instance, clique can be solved with only O( n) nondeterministic moves, and Pratt's algorithm [16] solves primality, which is not believed to be NP-complete, with O(n 2 ) nondeterministic moves. Motivated by the different amounts of nondeterminism apparently needed to solve problems in NP, Kintala and Fischer [9, 10, 11] defined limited nondeterminism classes within NP, including the classes we now call the fi hierarchy. The structural properties of the fi classes were studied further by ' Alvarez, Diaz and Toran [1, 6]. These classes arose yet again in the work of Papadimitriou and Yannakakis [15] on particular problems inside NP (e.g., quasigroup isomorphism can be solved with O(log 2 n) nondeterministic moves). Kintala and Fischer [11] defined P f(n) to be the class of languages accepted by a nondeterministic polynomial-time bounded Turing machine that makes at most c-ary nondeterministic moves (equivalently, O(f(n)) binary nondeterministic moves) on inputs of length n. Being mostly interested in polylogarithmic amounts of nondeterminism, they defined Diaz and Toran [6] wrote fi f(n) to denote Kintala and Fischer's P f(n) and fi k to denote PL k . Papadimitriou and Yannakakis [15] wrote NP[f(n)] to denote P f(n) . (Their work is surveyed in [7].) We will adopt the NP[f(n)] notation of Papadimitriou and Yannakakis, as well as the fi k notation of Diaz and Toran. To reiterate: Definition 1.1. ffl A language L belongs to NP[f(n)] if there exists a polynomial-time bounded nondeterministic Turing machine that accepts L and makes O(f(n)) nondeterministic choices on inputs of length n. (Note: NP[f(n)] ' DTIME(2 O(f(n)) ):) ffl The fi hierarchy consists of y Yale University, Dept. of Computer Science, P.O. Box 208285, New Haven, CT 06520-8285. email: [email protected]. Research supported in part by the United States National Science Foundation under grant CCR-8958528 and by the Netherlands Organization for Scientific Research (NWO) under Visitors Grant B 62-403. z University of Kentucky, Dept. of Computer Science, Lexington, KY 40506. email: [email protected]. Research supported in part by the National Science Foundation under grant CCR-9315354 Kintala and Fischer [11] constructed oracles that make the fi hierarchy collapse to any desired level. That is, there is an oracle relative to which and, for every k - 1, there is an oracle relative to which Oracles can also make the polynomial hierarchy and the Boolean hierarchy collapse to any desired level [12, 4]. The polynomial and Boolean hierarchies have a very nice property: collapses translate upward. I.e., if the kth and 1)st levels are equal, then all levels are contained in the kth [5, 4]. This is also reflected in the non-determinism hierarchy, now known as the b hierarchy, studied by Buss and Goldsmith [3]. The classes in the b hierarchy are defined by two parameters: the exponent of the polynomial time bound (ignoring log factors), and the constant factor for k log n bits of nondeterminism. This hierarchy exhibits upward collapse for both time and k. All attempts to prove an analogous translational property for the fi hierarchy have failed. In fact the obvious technique extends a collapse by only a constant factor in the number of nondeterministic bits, giving one of the aforementioned upward collapses of the b hierarchy. Hemachandra and Jha [8] attempted to explain this failure by examining the tally sets in the fi hierarchy. For each k, they constructed an oracle that makes We find this explanation unsatisfactory because it considers only tally sets. The known behavior of relativized fi hierarchies is that fi A 1 and that fi A A. A collapse is a statement of the form fi A . A separation is a statement of the form fi A . A closure is a statement of the form . A nonclosure is a statement of the form fi A . A requirement is a collapse, separation, closure, or nonclosure. We call a set of requirements consistent if it is consistent with the known behavior of relativized fi hierarchies, as stated above, and the standard axioms for Given a set S of requirements let X be the union of [0; 1] and all intervals [i; j] such that and the collapse fi A j belongs to S. It is easy to see that S is consistent iff the following conditions hold for all a and b such that [a; or [b; a] the separation fi A a 6= fi A b does not belong to S; (2) the closure a and the nonclosure fi A b do not both belong to S. For any consistent set of requirements, we construct an oracle A such that the fi hierarchy relative to A satisfies them. For example, for each k - 0, there is an oracle that makes fi Another oracle makes fi We can also make distinct levels in the hierarchy be closed under complementation or not, as long as this is consistent with the collapses (if fi A then we cannot have i and fi A We prove two initial results for every k: ffl There is an oracle that makes the first k levels of the fi hierarchy coincide, but makes the remaining levels all distinct (Theorem 2.3). ffl There is an oracle that makes the first k levels of the fi hierarchy coincide, the 1)st level different from the kth, and the remaining levels all equal (Theorem 2.5). The techniques from these two constructions can be combined to get any consistent finite or infinite set of collapses and separations. To collapse fi k into fi j (for DOWNWARD SEPARATION FAILS FOR fi HIERARCHY 3 k ? j), we code a complete set for fi k into fi j . The same coding techniques can also code fi i into fi i , for any i's we wish, as long as this doesn't violate any collapses. (If fi k is collapsed to fi j , then either both or neither will be closed under complement.) Finally, in Section 3, we extend our results to fi r for real r - 0. One theme in complexity theory is to ask whether contains any easy sets (assuming P 6= NP). The answer to the question above depends on the definition of "easy." Ladner [14] showed that if P 6= NP then contains an incomplete set. On the other hand, there are oracles relative to which P 6= NP, but contains (a) no tally sets [13] or (b) no sets in co-NP [2]. It is unknown whether the assumption implies that contains a set in DTIME(n polylog ); a positive answer would improve many constructions in the literature. As a step toward understanding that question, we construct an oracle relative to which P 6= NP but contains no set in the fi hierarchy (Corollary 2.2). 2. Limited Nondeterminism Hierarchies. The construction below gives almost all the techniques used in subsequent theorems. Theorem 2.1. Let g 0 and g 1 be polynomial-time computable monotone increasing functions such that log n 2 o(g 1 (n)) and g 0 (n) 2 n O(1) . If g 0 (n O(1) there exists an oracle A such that P (and in fact there is a tally set in (NP[g 1 (n)]) A \Gamma (NP[g 0 (n)]) A ). Proof. Let C accepts x within s steps with oracle A, making at most g 0 (jxj) nondeterministic choicesg. Then C A is - p m -complete for for every A. Let p(n) be the polynomial time bound for some NP[g 0 (n)] oracle Turing machine recognizing C () . Let D A]g. Note that D A 2 NP[g 1 (n)] A . The construction consists of coding C A into A in a polynomial-time recoverable manner, making (NP[g 0 (n)]) A ' P A , while diagonalizing, i.e, guaranteeing that no machine recognizes the set D A , so P A 6= (NP[g 1 (n)]) A . At the end of the construction, we will have We refer to all strings beginning with 1 as coding strings. We refer to all strings beginning with 0 as diagonalizing strings. Assume that P () is enumerated by Turing machines P (), P () runs in time bounded by n i for all i, and for sufficiently large n. The construction proceeds in stages. At the end of stage s, A is decided for all strings of length up to n s , and D A is extended so that P A s does not recognize D A . The stage consists of one diagonalization, which determines n s , and continued encoding of C A . At stage s, choose n ? n s\Gamma1 such that n is a power of 2, the running time of P () s on inputs of length n is at most n s , and n satisfies an inequality to be specified below that is to be specified below that is true for almost all n. Let . The value of depends only on oracle strings of length coding strings of length - A up to length ' \Gamma 1. In order to diagonalize, P A s (x) must be calculated. But that computation may query coding strings that code computations of C A that are not yet decided, because those computations in turn rely on strings for which A is not yet decided. Those strings in turn may depend on other coding strings. Any diagonalizing strings that do not 4 R. BEIGEL AND J. GOLDSMITH already belong to A and are queried by P A s (x), or by the computation corresponding to a coding string that P A s (x) queries, or in the computation corresponding to a coding string that one of those computations queries, or so on, are restrained from A. We claim that there are more potential witnesses for x to be in D A than there are possible queries in such a cascade of queries, so deciding P A s (x) does not restrict our decision about D A (x). Because of the encoding of C A , a coding string z codes a computation that depends only on strings of length bounded by jzj. C A (w) directly depends on at most p(jwj)2 g0 (jwj) of these shorter strings. A computation of P A s (x) may query no more than n s strings, each of length bounded by n s . Each of these strings may code a computation on a string of length at most n s=2 . Each of these computations depends on at most p(n s=2 )2 g0 (n s=2 ) strings, each of which depends on at most p(n s=4 )2 g0 (n s=4 ) strings, etc. This recursion can be cut off at strings of length ' \Gamma 1, because A is already determined up to that length. The total number of queries needed to decide P A s (x) is bounded by n s times the product of all the terms above of the form p(n s=2 i There are at most log log n s \Gamma log log log log n s \Gamma log log log s such terms, and each of them is bounded by p(n s )2 g0 (n s Therefore the total number of queries on which P A s (x) depends is n s 2 O(log(s)g0 (n s )) , which is less than 2 g1 (n) for sufficiently large n. (The inequality that n must satisfy is n s (p(n s )2 g0 (n s ) Thus there remains an unrestrained diagonalizing string of length ', which we put into A if P A s (x) rejects x. That is, we set D A s (x), adding a string 0xy to A if necessary. Thus, for each s, we can guarantee that P A s does not accept D A , so A , this shows that (NP[g 1 (n)]) A 6= P A . Since is complete for (NP[g 0 (n)]) A , this shows that (NP[g 0 (n)]) A ' P A . We complete stage s by letting n finishing the coding of any C A (w) that was begun or queried in this stage. The preceding theorem is tight because if (even if we restrict to binary nondeterministic moves) via a relativizable proof. (Pre- viously, Sanchis [17] had observed that if Because the classes are separated by tally sets, we also separate the exponential-time versions of these classes (see [8] for elaboration of this). Corollary 2.2. There is an oracle relative to which Proof. Let and in the previous theorem. Then for all k, fi A Theorem 2.3. Let g(\Delta; \Delta) be a polynomial-time computable, monotone increasing (in both variables) function with log there exists an oracle A such that (and in fact there is a tally set in (NP[g(n; 2i for each i). (In this theorem, we ignore the relationship between NP[g(n; NP[g(n; 2i)]. We will take that up in the next theorem.) The only difference between this construction and the previous one is that there are infinitely many diag- onalizations going on. At stage we guarantee that the eth machine for (NP[g(n; 2i)]) A does not accept the diagonal set D A . Thus, . The counting argument for this construction is identical to that in the proof of Theorem 2.1. DOWNWARD SEPARATION FAILS FOR fi HIERARCHY 5 Corollary 2.4. For any k, there is an oracle relative to which Proof. Let g(n; in the preceding theorem. Theorem 2.5. Let g 0 and g 1 be polynomial-time computable monotone increasing functions such that log n 2 o(g 1 (n)) and g 0 (n) 2 n O(1) . If g 0 (n O(1) there exists an oracle A such that (and in fact there is a tally set in (NP[g 1 (n)]) A \Gamma (NP[g 0 (n)]) A ). Sketch. In this construction, we do two encodings and one diagonalization. In addition to coding C A into P, we also code E A , a generic - p m -complete set for PSPACE, into A. accepts x using at most s tape squares with oracle Ag, where we also count the space used on the oracle tape.) At the end of the construction, we have (If one prefers binary oracles, one may code 0, 1, and 2 as 00, 01, and 10.) When we are doing a diagonalization to make P A s (x) 6= D A (x), if a coding string for C A k (w) is queried, we proceed as before; if a coding string for E A (w) is queried, where jwj - jxj, then we simply restrain that coding string from the oracle. This will not restrain all the coding strings for E A (w), since there are 2 g1 (jwj) coding strings for E A (w); if is the upper bound on the total number of queries generated by the computation of P A as in the proof of Theorem 2.1. Therefore, restraining any such coding strings queried in the computation of P A s (x) or in its cascade of queries can not restrain all such coding strings, and thus can not decide E A (w). At the end of each stage, we complete all codings begun or queried in that stage, so that it will not be changed in any subsequent stage. Corollary 2.6. For every k, there is an oracle relative to which With only a slight modification of this technique, we get far more bizarre collapses. Theorem 2.7. Let g(\Delta; \Delta) be a polynomial-time computable, monotone increasing (in both variables) function with log n 2 o(g(n; i)) and g(n; i) 2 n O(1) for all i - 1. there exists an oracle A such that for all i - 0 (and in fact there is a tally set in (NP[g(n; 2i for each i). We include the full proof of this result, although it uses techniques mentioned before, since this shows how all the pieces fit together. Proof. Let C A accepts x within s steps with oracle A, making at most g(jxj; i) nondeterministic choicesg. Then C A m -complete for (NP[g(n; i)]) A for any A. Let p(n; i) be the nondeterministic time bound for some 6 R. BEIGEL AND J. GOLDSMITH Turing machine recognizing C () . Without loss of generality, assume that for all i and almost all n, Let D A A]g. For convenience, define g(n; The construction consists of coding C A 2i into (NP[g(n; 2i \Gamma 1)]) A , for each i - 0, so (NP[g(n; 0)]) A ' P A and (NP[g(n; 2i diagonalizing, i.e, guaranteeing that no (NP[g(n; 2i)]) A machine recognizes the set D A 2i+1 , for any i, so (NP[g(n; 2i + 1)]) A 6' (NP[g(n; 2i)]) A for any i. At the end of the construction, we will have x 2 C A Assume that (NP[g(n; i)]) A is enumerated by oracle NTMs M runs in time bounded by n e for sufficiently large n. The construction proceeds in stages. Stage consists of some encodings and one diagonalization, which determines n s . At the end of stage s, A is decided for all strings of length - n s (and some further coding strings), and A has been extended so that M A e;2i does not recognize D A 2i+1 . At stage s, let he; ii = s, and then choose n ? n s\Gamma1 such that n is a power of 2, runs in time bounded by n e on inputs of length n, and n satisfies an inequality to be specified below that is true for almost all n. Let n . The value of D A depends only on strings of length 1). Do all coding involving witnesses of length less than ', and then freeze A through length ' \Gamma 1. As before, in order to diagonalize, we will need to calculate M A e;2i (x), which may generate a cascade of queries. Any diagonalizing strings that do not already belong to A and are queried in this cascade are restrained from A. But coding strings may be queried as well. (Because we are coding nondeterministically, coding strings can be thought of as potential witnesses to membership.) If M A queries a potential witness that w 2 C A 2j (w) has not yet been decided, that potential witness is restrained from A. If 2j - 2i and C A 2j (w) has not yet been decided, then we compute C A 2j (w) recursively. We will show below that the number of queries generated by such a cascade of queries is smaller than both of the following bounds: (1) the number of potential witnesses for w 2 C A 2j , (2) the number of potential . In fact, bound (1) implies bound (2) as follows. The number of witnesses for D A 2i+1 (x) is 2 g(jxj;2i+1) , and the number of witnesses for C A 2j (w) is . If a witness of C A 2j (w) is restrained, then jwj - jxj and 2j ? 2I . Thus by monotonicity of g, g(jwj; Thus, restraining potential witnesses as described does not impede any encodings or restrict our decision about D 2i+1 (x), or those C A 2j (w) for which we restrict coding strings. (We don't have to worry about what happens to potential witnesses for 2j at a later stage, because any affected codings, i.e., C A 2j (w), will be completed at this stage; later diagonalizations will not affect them.) Now we show that there are more potential witnesses for x 2 D A 2i+1 than there are possible queries in such a cascade of queries. Because of how we encode C A k , a coding string z codes a computation that depends on strings of length bounded by jzj. For (w) depends on at most p(jwj; k)2 these shorter strings. e;2i (x) has at most 2 g(n;2i) computations, and each of those computations may query no more than n e strings, each of length bounded by n e . Each such string may code a computation C A e , but we only need to expand that compu- DOWNWARD SEPARATION FAILS FOR fi HIERARCHY 7 tation if 2j - 2i. Each of these computations depends on at most p(n e=2 ; 2i)2 g(n e=2 ;2i) strings, each of which depends on at most p(n e=4 ; 2i)2 g(n e=4 ;2i) strings, etc. As be- fore, the total number of queries needed to decide M A e;i (x) is bounded by the product of log e - log s terms, each of which is 2 o(g(n;2i+1)) . Therefore the total number of queries on which M A depends is 2 o(g(n;2i+1)) , which is less than 2 g(n;2i+1) for sufficiently large n. Thus there remains an unrestrained diagonalizing string of length ', which we put into A if and only if M A rejects x. That is, we set D A adding a string 0xy to A if necessary. Thus, for each s, we can guarantee that M A does not accept D A this shows that (NP[g(n; 2i Since C A 2i is complete for (NP[g(n; 2i)]) A , our encoding guarantees that Corollary 2.8. There is an oracle relative to which, for each k, fi fi 2k+2 . Corollary 2.9. For any consistent pattern of collapses and separations of the k 's, there is an oracle relative to which that pattern holds. Notice that if the set of collapses is not recursive, then the oracle will also be non-recursive. In addition to collapsing or separating fi j and fi k , we can code co-fi k into fi k - or separate the two. This involves some additional argument. Theorem 2.10. Let g(\Delta; \Delta) be a polynomial-time computable, monotone increasing (in both variables) function with log n 2 o(g(n; i)) and g(n; i) 2 n O(1) for all there exists an oracle A such that P (and in fact there are tally sets in (NP[g(n; 2i A and in (NP[g(n; 2i for each i). Sketch. For convenience, we will separate (NP[g(n; 2i+1)]) A from (co-NP[g(n; 2i rather than (NP[g(n; 2i 2)]) A from (co-NP[g(n; 2i given the other re- quirements, this is equivalent. In order to separate (NP[g(n; 2i 1)]) A from we use the set D A 2i+1 in (NP[g(n; 2i that D A . Most of this construction is identical to that of Theorem 2.7, except that we interleave an extra diagonalization into the construc- tion; the codings and diagonalizations are analogous to earlier constructions, and the counting argument is identical. We code complete sets C A 2i for (NP[g(n; 2i)]) A into (NP[g(n; 2i \Gamma 1)]) A , and diagonalize so that no (NP[g(n; 2i)]) A machine recognizes D A (Thus D A does double duty: during even stages, it diagonalizes against during odd stages, against (co-NP[g(n; 2i To guarantee that D A 2i+1 is not in (co-NP[g(n; 2i , we make sure that, for each e, the e th machine for (NP[g(n; 2i does not recognize D A . This holds if and only if there is some x such that D A e;2i+1 (x). This diagonalization differs from earlier ones only when M A queries a witness for x 2 D A 2i+1 . As before, if M A queries a coding string for some computation of C A 2j (w) where then we can safely restrain the coding string. (If may exclude w from D A 2i+1 , but that doesn't matter. As long as D A e;2i+1 (x), we don't care what happens to D A 2i+1 for other strings 8 R. BEIGEL AND J. GOLDSMITH of lengths between n s\Gamma1 and n s , where then we retrace the computation, as before. If M A queries a witness for x 2 D A 2i+1 , we first restrain all such wit- nesses, and continue. If this leads to a rejecting computation of M A 2i+1 (x), and the diagonalization is successful. If it leads to an accepting computation, we preserve the lexicographically least accepting path for that computation, and all of its cascade of queries. As before, the computation of restrains at most 2 o(g(n;2i+1)) strings, so this will not restrain all the witnesses for x 2 D A 2i+1 . Thus we can find an unrestrained witness and add it to A, so D A e;2i+1 (x), as desired. Therefore, this additional set of diagonalization requirements can be interleaved with the previously-described diagonalizations and collapses. Theorem 2.11. Let g(\Delta; \Delta) be a polynomial-time computable, monotone increasing (in both variables) function with log n 2 o(g(n; i)) and g(n; i) 2 n O(1) for all there exists an oracle A such that P (and in fact there are tally sets in (NP[g(n; 2i for each i). Sketch. As before, we construct A so that no (NP[g(n; 2i)]) A machine recognizes the set D A 2i+1 , and so that C A In addition, in order to make as follows: For each i, let N A i be an (NP[g(n; i)]) A machine recognizing C A i in nondeterministic time bounded by p(n; i) (regardless of the oracle). By the form of the encoding, query any of its own coding strings. If a witness string for C A is queried in the course of a diagonalization (NP[g(n; 2i)]) A 6= (NP[g(n; 2i then we can retrace the computation. If 2j we can restrain the queried witness string (for jxj sufficiently large) without deciding C A 2j+1 (x), by the same counting argument as in previous proofs. Thus, we can add this extra encoding, without interfering with the other collapses and codings. This gives us the following stronger version of Hemachandra and Jha's oracle [8]. Corollary 2.12. For each k, there is an oracle relative to which for all j, (and the separations are witnessed by tally sets). Combining the results (and techniques) of Theorems 2.7, 2.10, and 2.11, we get the following very strong result. Corollary 2.13. For any consistent set of requirements, there is an oracle relative to which the fi hierarchy satisfies those requirements. In constructing such an oracle, one must be careful in closing classes under com- plement. In particular, if we close one class under complement, and separate another from its complement, we cannot then make the two classes equal. Corollary 2.9 implies that there are uncountably many different patterns of collapse that can be realized in relativized worlds. If the set of requirements is recursive, then the oracle can be made recursive, but certainly some of those patterns are realized by only nonrecursive oracles. DOWNWARD SEPARATION FAILS FOR fi HIERARCHY 9 3. Dense fi Hierarchies. Previously we considered fi r only when r is a natural number. But the class (NP[log r n]) A is meaningful whenever r is a nonnegative real number (regardless of whether r is computable). Even when we allow r to be real, we can make the fi hierarchy obey any consistent set of requirements. For example, we can make the fi hierarchy look like a Cantor set. Theorem 3.1. Let X be any subset of [1; 1). There exists an oracle A such that, for all s; t - 1, fi A t if and only if [s; t] ' X. Note that there may be uncountably many distinct fi t s. Because there are two rationals between any two reals, we need only separate the distinct fi q s where q is rational. Proof. Without loss of generality, assume that X is a union of intervals, each containing more than one point. Every interval in X contains a rational point; therefore contains countably many intervals. We will satisfy the following requirements for each maximal interval in X , depending on its type: log log n] log log n] In addition, for each rational number q in (1; NP[log q n= log log n] 6= NP[log q n] 6= NP[log q n log log n]: then we make P 6= NP[log n log log n] as well. The construction is a slight modification of that in the proof of Theorem 2.7. We perform the diagonalizations in some well-founded order, while maintaining the codings as we go along. The only significant difference here is that the diagonalizations are not performed in increasing order. Suppose that at some stage we are making ae NP[b(n)] and a coding string for some NP[c(n)] computation is queried; we restrain that string if and only if (9n)[c(n) ? b(n)] if and only if (8n)[c(n) - b(n)]. The counting argument is the same as before. Note: we could also close each distinct fi r under complement or not, as we wish, in the theorem above. 4. Open Problems. The class fi k is contained in NP " DTIME(2 log k n ). Our work shows that there is no relativizing proof that We would like to know whether are there any easy languages in NP \Gamma P? The best we can show is that if well-behaved function f , then Is there an oracle relative to which this is the best possible translation of the collapse? Does Acknowledgments . We are grateful to Leen Torenvliet, Andrew Klapper, and Martin Kummer for helpful discussions, and Andrew Klapper, Bill Gasarch, and Martin Kummer for careful proofreading of earlier drafts. --R "Complexity Classes With Complete Problems Between P and NP-Complete," Relativizations of the P "Nondeterminism within P," The Boolean hierarchy I: structural properties. Classes of bounded nondeterminism. Limited nondeterminism SIGACT News Defying upward and downward separation. Computations with a restricted number of nondeterministic steps. Computations with a restricted number of nondeterministic steps. Refining nondeterminism in relativized polynomial-time bounded computations Relativized polynomial hierarchies having exactly k levels. Sparse sets in NP On the structure of polynomial time reducibility. On limited nondeterminism and the complexity of the V-C dimension Every prime has a succinct certificate. Constructing language instances based on partial information. --TR

oracles;structural complexity theory;hierarchies;limited nondeterminism

290287

A rejection technique for sampling from log-concave multivariate distributions.

Different universal methods (also called automatic or black-box methods) have been suggested for sampling form univariate log-concave distributions. The descriptioon of a suitable universal generator for multivariate distributions in arbitrary dimensions has not been published up to now. The new algorithm is based on the method of transformed density rejection. To construct a hat function for the rejection algorithm the multivariate density is transformed by a proper transformation T into a concave function (in the case of log-concave density log(x).) Then it is possible to construct a dominating function by taking the minimum of serveral tangent hyperplanes that are transformed back by T-1 into the original scale. The domains of different pieces of the hat function are polyhedra in the multivariate case. Although this method can be shown to work, it is too slow and complicated in higher dimensions. In this article we split the Rn into simple cones. The hat function is constructed piecewise on each of the cones by tangent hyperplanes. The resulting function is no longer continuous and the rejection constant is bounded from below but the setup and the generation remains quite fast in higher dimensions; for example, 8. The article describes the details of how this main idea can be used to construct algorithm TDRMV that generates random tuples from a multivariate log-concave distribution with a computable density. Although the developed algorithm is not a real black box method it is adjustable for a large class of log-concave densities.

Introduction For the univariate case there is a large literature on generation methods for standard distributions (see e.g. [Dev86] and [Dag88]) and in the last years some papers appeared on universal (or black-box) methods (see [Dev86, chapter VII], [GW92], [Ahr95], [H-or95a], [HD94] and [ES97]); these are algorithms that can generate random variates from a large family as long as some information (typ- ically the mode and the density of the specific distribution) are available. For the generation of variates from bivariate and multivariate distributions papers are rare. Well known and discussed are only the generation of the multi-normal and of the Wishart distribution (see e.g. [Dev86] and [Dag88]). Several approaches to the problem of generating multivariate random tuples exist, but these have some disadvantages: ffl The multivariate extension of the ratio of uniforms methods as in [SV87] or [WGS91]. This method can be reformulated as rejection from a small family of table-mountain shaped multivariate distributions. This point of view is not included in these two papers but it is useful as it clarifies the question why the acceptance probability becomes poor for high correla- tion. This disadvantage of the method is already mentioned in [WGS91]. The practical problem how to obtain the necessary multivariate rectangle enclosing the region of acceptance for the ratio of uniforms method is not discussed in [SV87] nor in [WGS91] and seems to be difficult for most distributions. ffl The conditional distribution method. It requires the knowledge of and the ability to sample from the marginal and the conditional distributions (see [Dev86, chapter XI.1.2]). ffl The decomposition and rejection method. A majorizing function (also called suggested for the multivariate rejection method is the product of the marginal densities (in [Dag88]). It is not clear at all how to obtain the necessary rejection constant ff. ffl Development of new classes of multivariate distributions, which are easy to generate. It is only necessary (and possible) to specify the marginal distribution and the degree of dependence measured by some correlation coefficient (see the monograph [Joh87]). This idea seems to be attractive for most simulation practitioners interested in multivariate distributions but it is no help if variates from a distribution with given density should be generated. Recently Devroye [Dev97] has developed algorithms for ortho-unimodal densities. But this paper leaves the generation of log-convave distributions as open problem. ffl Sweep-plane methods for log-concave (and T-concave) distributions are described recently in [H-or95b] for bivariate case and in [LH98] for the multivariate case. These algorithms use the idea of transformed density rejection which is presented in a first form in [Dev86, chapter VII.2.4] and with a different set-up in [GW92]. To our knowledge these two algorithms are the only universal algorithms in the literature for multivariate distributions with given densities. (In [Dev86, chapter XI.1.3] it is even stressed that no general inequalities for multivariate densities are available, a fact which makes the design of black-box algorithms, that are similar to those developed in [Dev86] for the univariate case, impossible.) Although the algorithm in [LH98] works, it is very slow, since the domain of the density f is decomposed in polyhedra. This is due to the construction of the hat function, where we take the pointwise minimum of tangent hyperplanes. In this paper we again use transformed density rejection and the sweep-plane technique to derive a much more efficient algorithm. The main idea is to decompose the domain of the density in cones first and then compute tangent hyperplanes in this cones. The resulting hat function is not continuous any more and the rejection constant is bounded from below, but the setup as well as the sampling from the hat function is much faster than in the original algorithm. Section 2 explains the method and gives all necessary mathematical formu- lae. Section 3 provides all details of the algorithm. Section 4 discusses how to improve and extend the main idea of the algorithm (e.g. to T-concave distribu- tions, bounded domain) and section 5 reports the computational experience we have had with the new algorithm. 2 The method 2.1 Transformed density rejection Density. We are given a multivariate distribution with differentiable density function To simplify the development of our method we assume In x4 we extend the algorithm so that these requirements can be dropped. Transformation. To design an universal algorithm utilizing the rejection method it is necessary to find an automatic way to construct a hat function for a given density. Transformed density rejection introduced under a different name in [GW92] and generalized in [H-or95a] is based on the idea that the density f is transformed by a monotone T (e.g. T in such a way that (see [H-or95a]): concave (we then say "f is T -concave"); differentiable and T 0 (x) ? 0, which implies T \Gamma1 exists; and (T4) the volume under the hat is finite. Hat. It is then easy to construct a hat ~ h(x) for ~ f(x) as the minimum of N tangents. Since ~ f(x) is concave we clearly have ~ Transforming ~ h(x) back into the original scale we get majorizing function or hat for f , i.e. with f(x) - h(x). Figure 1 illustrates the situation for the univariate case by means of the normal distribution and the transformation T log(x). The left hand side shows the transformed Figure 1: hat function for univariate normal density density with three tangents. The right hand side shows the density function with the resulting hat. (The dashed lines are simple lower bounds for the density called squeezes in random variate generation. Their use reduces the number of evaluations of f . Especially if the number of touching points is large and the evaluation of f is slow the acceleration gained by the squeezes can be enormous.) Rejection. The basic form of the multivariate rejection method is given by algorithm Rejection(). Algorithm 1 Rejection() 1: Set-up: Construct a hat-function h(x). 2: Generate a random tuple proportional to h(X) and a uniform random number U . 3: If Uh(X) - f(X) return X else go to 2. The main idea of this paper is to extend transformed density rejection as described in [H-or95a] to the multivariate case. 2.2 Construction of a hat function Tangents. Let p i be points in D ' R n . In the multivariate case the tangents of the transformed density ~ f(x) at p i are the hyperplanes given by where h\Delta; \Deltai denotes the scalar product. Polyhedra. In [LH98] a hat function h(x) is constructed by the pointwise minimum of these tangents. We have The domains in which a particular tangent ' i (x) determines the hat function are simple convex polyhedra P i , which may be bounded or not (for details about convex polyhedra see [Gr-u67, Zie95]). Then a sweep-plane technique for generating random tuples in such a polyhedron with density proportional to To avoid lots of indices we write p, '(x) and P without the index i if there is no risk of confusion. A sweep-plane algorithm. Let r ~ kr ~ if r ~ choose any g with denotes the 2-norm.) For a given x let We denote the hyperplane perpendicular to g through x by fy and its intersection with the polytope P with depends on x only; thus we write F (x), if there is no risk of confusion.) Q(x) again is a convex simple polyhedra. Now we can move this sweep-plane F (x) through the domain P by varying x. Figure 2 illustrates the situation. As can easily be seen from (2), (4) and (5), T \Gamma1 ('(x)) is constant on Q(x) for every x. Let Then the hat function in P is given by where again We find for the marginal density function of the hat Z Figure 2: sweep-plane F (x) where integration is done over F (x). A(x) denotes the (n \Gamma 1)-dimensional volume of Q(x). It exists if and only if Q(x) is bounded. To compute A(x) let v j denote the vertices of P and v assume that the polyhedron P is simple. Then let t v j n be the n nonzero vectors in the directions of the edges of P originated from v j , i.e. for each k and every by modifying the method in [Law91] we find a The coefficients are given by Y and a Notice that b (x) equations (9) and (10) does not hold if P is not simple. For details see [LH98]. The generation from h g is not easy in general. But for log-concave or -concave (see x4.8) densities f(x), h g again is log-concave ([Pr'e73]) and T c - concave ([LH98]), respectively. Generate random tuples. For sampling from the "hat distribution" we first need the volume below the hat in all the polyhedra P i and in the domain D. We then choose one of these polytopes randomly with density proportional to their volumes. By means of a proper univariate random number we sample from marginal distribution hj g and get a intersection Q(x) of P . At last we have to sample from a uniform distribution on Q(x). It can be shown (see [LH98]) that the algorithm works if (1) the polyhedra P i are simple (see above), (2) there exists a unique maximum of ' i (x) in P i (then ff \Gamma fi x is decreasing and thus the volume below the hat is finite in unbounded polyhedra), and is non-constant on every edge of P i (otherwise hg; t v j and an edge t i and thus a Adaptive rejection sampling. It is very hard to find optimal points for constructing these tangents ' i (x). Thus these points must be chosen by adaptive rejection sampling (see [GW92]). Adapted to our situation it works in the following way: We start with the vertices of a regular simplex and add a new construction point whenever a point is rejected until the maximum number N of tangents is reached. The points of contact are thus chosen by a stochastic algorithm and it is clear that the multivariate density of the distribution of the next point for a new tangent is proportional to h(x) \Gamma f(x). Hence with tending towards infinity the acceptance probability for a hat constructed in such a way converges to 1 with probability 1. It is not difficult to show that the expected volume below the hat is 1 +O(N \Gamma2=n ). Problems. Using this method we run into several problems. We have to compute the polyhedra every time we add a point. What must be done, if the marginal distribution (8) does not exist in the initial (usually not bounded) polyhedra P i , or if the volume below the hat is infinite (Q i (x) not bounded, ff \Gamma fi x not decreasing)? Moreover the polyhedra P i typically have many vertices. Therefore the algorithm is slow and hard to implement because of the following effects. \Gamma The computation of the polyhedra (setup) is very expensive. \Gamma The marginal density (8) is expensive to compute. Since it is different for every polyhedron P i (and for every density function f ), we have to use a slow black box method (e.g. [GW92, H-or95a]) for sampling from the marginal distribution even in the case of log-concave densities. \Gamma Q(x) is not a simplex. Thus we have to use the (slow) recursive sweep- plane algorithm as described in [LH98] for sampling from the uniform distribution over a (simple) polytope. 2.3 Simple cones A better idea is to choose the polyhedra first as simple as possible, i.e. we choose cones. (We describe in x2.4 how to get such cones.) A simple cone C (with its vertex in the origin) is an unbounded subset spanned by n linearly independent vectors: In opposition to the procedure described above we now have to choose a proper point p in this cone C for constructing a tangent. In the whole cone the hat h is then given by this tangent. The method itself remains the same. Obviously the hat function is not continuous any more (because we first define a decomposition of the domain and then compute the hat function over the different parts. It cannot be made continuous by taking the pointwise minimum of the tangents, since otherwise we cannot compute the marginal density h g by equation (8)). Moreover we have to choose one touching point in each part. These disadvantages are negligible compared to the enormous speedup of the setup and of the generation of random tuples with respect to this hat function. Marginal density. The intersection Q(x) of the sweep plane F (x) with the cone C is bounded if and only if F (x) cuts each of the sets f-t x ? 0, i.e. if and only if hg; -t hence if and only if We find for the volume A(x) in (9) of the intersection Q(x) ae a x where (again) a Y Notice that A(x) does not exist if condition (13) is violated, whereas the right hand side in (14) is defined If the marginal density exists, i.e. (13) holds, then by (8) and (6) it is given by Volume. The volume below the hat function in a cone C is given by Z 1h g Z 1a x Notice that g and thus a, ff and fi depend on the choice of p. Choosing an arbitrary p may result in a very large volume below the hat and thus in a very poor rejection constant. Intersection of sweep-plane. Notice that the intersection Q(x) is always a (n \Gamma 1)-simplex if condition (13) holds. Thus we can use the algorithm in [Dev86] for sampling from uniform distribution on Q(x). The vertices of Q(x) in R n are given by uniformly [0; 1] random variates and U We sort these variates such that U 0 - Un . Then we get a random point in Q(x) by (see [Dev86, theorems XI.2.5 and V.2.1]) The choice of p. One of the main difficulties of the new approach is the choice of the touching point p. In opposite to the first approach where the polyhedron is build around the touching point, we now have to find such a point so that holds. Moreover the volume below the hat function over the cone should be as small as possible. Searching for such a touching point in the whole cone C or in domain D (the touching point needs not to be in C) with techniques for multidimensional minimization is not very practicable. Firstly the evaluation of the the volume HC in (17) for a given point p is expensive and its gradient with respect to p is not given. Secondly the domain of HC is given by the set of points where (13) holds. Instead we suggest to choose a point in the center of C for a proper touching point for our hat. Let - be the barycenter of the spanning vectors. Let a(s), ff(s) and fi(s) denote the corresponding parameters in (16) for t. Then we choose by minimizing the function Z The domain DA of this function is given by all points, where kr ~ where A(x) exists, i.e. where condition (13). It can easily be seen, that DA is an open subset of (0; 1). To minimize j we can use standard methods, e.g. Brent's algorithm (see e.g. [FMM77]). The main problem is to find DA . Although ~ f(x) is concave by assumption, it is possible for a particular cone C that DA is a strict subset of (0; 1) or even the empty set. Moreover it might not be connected. In general only the following holds: Let (a; b) be a component of DA 6= ;. If If f 2 C 1 , i.e. the gradient of f is continuous, then lim s&a Roughly spoken, j is a U-shaped function on (a; b). An essential part of the minimization is initial bracketing of the minimum, i.e. finding three points s (a; b), such that j(s 1 This is necessary since the function term of j in (20) is also defined for some s 62 DA (e.g. s ! 0). Using Brent's algorithm without initial bracketing may (and occasionally does) result in e.g. a negative s. Bracketing can be done by (1) search for a s 1 2 DA , and (2) use property (21) and move towards a and b, respectively, to find an s 0 and an s 2 . (It is obvious that we only find a local minimum of j by this procedure. But in all the distributions we have tested, there is just one local minimum which therefore is the global one.) For the special case where hg(s); - ti does not depend on s (e.g. for all multivariate normal distributions) DA either is (0; 1) or the empty set. It is then possible to make similar considerations like that in [H-or95a, theorem 2.1] for the one dimensional case. Adapted to the multivariate case it would state, that for the optimal touching point p, f(p) is the same for every cone C. Condition violated. Notice that DA even may be the empty set, i.e. condition fails for all s 2 (0; 1). By the concavity of ~ f(x) we know, that construction point p. Furthermore hg; pi is bounded from below on every compact subset of the domain D of the density f . Therefore there always exists a partition into simple cones with proper touching points which satisfy (13), i.e. the domains DA are not empty for all cones C. We even can have 2.4 Triangulation For this new approach we need a partition of the R n into simple cones. We get such a partition by triangulation of the unit sphere S n\Gamma1 . Each cone C is then generated by a simplex \Delta ae S n\Gamma1 (triangle in S 2 , tetrahedron in S 3 , and so \Deltag (22) These simplices are uniquely determined by the vectors t their vertices. (They are the the convex hull of these vertices in S n\Gamma1 .) It does not matter that these cones are closed sets. The intersection of such cones might not be empty but has measure zero. For computing a in (15) we need the volumes of these simplices. To avoid DA being the empty set, some of the cones have to be skinny. Furthermore to get a good hat function, these simplices should have the same volume (if possible) and they should be "regular", i.e. the distances from the center to the vertices should be equal (or similar). Thus the triangulation should have the following properties: Recursive construction. are easy computable for all simplices. (C3) Edges of a simplex have equal length. Although it is not possible to get such a triangulation for n - 3 we suggest an algorithm which fulfils (C1) and (C2) and which "nearly" satisfies (C3). Initial cones. We get the initial simplices as the convex hull in S n\Gamma1 of the vectors en (23) where e i denotes the i-th unit vector in R n (i.e. a vector where the i-th component is 1 and all others are 1g. As can easily be seen the resulting partition of the R n is that of the arrangement of the hyperplanes Hence we have 2 n initial cones. Barycentric subdivision of edges. To get smaller cones we have to triangulate these simplices. Standard triangulations of simplices which are used for example in fixed-point computation (see e.g. [Tod76, Tod78]) are not appropriate for our purpose. The number of simplices increases too fast for each triangulation step. (In opposition to fixed point calculations, we have to keep all simplices with all their parameters in the memory of the computer.) Instead we use a barycentric subdivision of edges: Let t be the vertices of a simplex \Delta. Then use the following algorithm. (1) Find the longest edge (2) Let i.e. the barycenter of the edge projected to the sphere. (3) Get two smaller simplices: Replace vertex t i by t new for the first simplex and vertex t j by t new for the second one. We have After making k of such triangulation steps in all initial cones we have 2 n+k simplices. This triangulation is more flexible. Whenever we have a cone C, where D a is empty (or the algorithm does not find an s 2 D a ) we can split C and try again to find a proper touching point in both new cones. This can be done until we have found proper construction points for all cones of the partition (see end of x2.3). In practice this procedure stops, if too many cones are necessary. (The computer runs out of memory.) Notice that it is not a good idea to use barycentric subdivision of the whole simplex (instead of dividing the longest edge). This triangulation exhibits the inefficient behavior of creating long, skinny simplices (see remark in [Tod76]). "Oldest" edge. Finding the longest of the edges of a simplex is very expensive. An alternative approach is to use the "oldest" edge of a simplex. The idea is the following: (1) Enumerate the 2n vertices of the initial cones. (2) Whenever a new vertex is created by barycentric subdivision, it gets the next number. (3) Edges are indexed by the tuple (i; j) of the number of the incident vertices, such that i ! j. We choose the edge with the lowest index with respect to the lexicographic order (the "oldest" edge). This is just the pair of lowest indices of the vertices of the simplex. As can easily be seen, the "oldest" edge is (one of) the longest edge(s) for the first steps. Unluckily this does not hold for all simplices in following triangulation steps. (But it is at least not the shortest one.) Computational experiences with several normal distributions for some dimensions have show, that this idea speeds up the triangulation enormously but has very little effect on the rejection constant. Setup. The basic version of the setup algorithm is as follows: 1. Create initial cones. 2. Triangulate. 3. Find touching points p if possible (and necessary). 4. Triangulate every cone without proper touching point. 5. Goto 3 if cones without proper touching points exist, otherwise stop. 2.5 Problems Although this procedure works for our tested distributions, an adaptation might be necessary for a particular density function f . (1) The searching algorithm for a proper touching point in x2.3 can be im- proved. E.g. DA is either [0; 1) or the empty set if f is a normal distribution. (2) There is no criterion how many triangulation steps are necessary or usefull for an optimal rejection constant. Thus some tests with different numbers of trianglation steps should be made with density f (see also x5). (3) It is possible to triangulate each cone with a "bad" touching point. But besides the case where no proper touching point can be found, some touching points may lead to an enourmous volume below the hat function. So this case should also be excluded and the corresponding cones should be triangulated. A simple solution to this problem is that an upper bound Hmax for the volumes HC is provided. Each cone with HC ? Hmax has to be triangulated further. Such a bound can be found by some empirical tests with the given density f . Another way is to triangulate all initial cones first and then let Hmax be a multiple (e.g. 10) of the 90th percentile of the HC of all created cones. Problems might occur when the mode is on the boundary of the support (Then we set can be seen as a concave An example for such a situation is when f(x) is a normal density on a ball B and vanishes outside of B. In such a case there exists a cone C such that f- t: - ? 0g does not intersect suppf and the algorithm is in troubles. If C " we simply can remove this cone. Otherwise an expensive search for a proper touching point is necessary. Restrictions. The above observations - besides the fact that no automatic adaption is possible - are a drawback of the algorithm for its usage as black-box algorithm. Nevertheless the algorithm is suitable for a large class of log-concave densities and it is possible to include parameters into the code to adjust the algorithm for a given density easily. Of course some tests might be necessary. Besides, the algorithm does not produce wrong random points but simply does not work, if no "good" touching points can be found for some cones C. 2.6 Log-concave densities The transformation T is concave, we say f is log-concave. We have T thus we find for the marginal density function in (16) those of a gamma distribution with shape parameters n and fi. The volume below the hat for log-concave densities in a cone C is now given by Z 1a x To minimize this function it is best to use its logarithm: For the normal distribution with density proportional to we have ~ is the center of the cone C with 1. Thus we simply find by (6) Since a(s) does not depend on s we find for constant. But even for the normal distribution with an arbitrary covariance matrix, this function becomes much more complicated. 3 The algorithm The algorithm tdrmv() consists of two main parts: the construction of a hat function h(x) and the generation of random tuples X with density proportional to this hat function. The first one is done by the subroutine setup(), the second one by the routine sample(). Algorithm random tuple for given log-concave density Input: density f 1: call Construct a hat-function h(x) 2: repeat 3: X / call sample(). = Generate a random tuple X with density prop. to h(X). 4: Generate a uniform random number U . 5: until U \Delta h(X) - f(X). return X. To store h(x), we need a list of all cones C. For each of these cones we need several data which we store in the object cone. Notice that the variables p, g, ff, fi, a and HC depend on the choice the touching point p and thus on s. Some of the parameters are only necessary for the setup. object 1 cone parameter variable definition spanning vectors t center of cone - construction point location of p s (setup) sweep plane g see (4) marginal density ff; fi see (6) coefficient a see (15) determinant of vectors volume under hat HC ; H cum C see (17) and (28) Remark. To make the description of the algorithm more readable, some standard techniques are not given in details. 3.1 The routine setup() consists of three parts: (H1) setup initial cones, (H2) triangulation of the initial cones and (H3) calculation of parameters. (H1) is simple (see x2.4). (H2) is done by subroutine split(). The main problem in (H3) is how to find the parameter s (i.e. a proper construction point). This is done by subroutine find(). Minimizing (29) is very expensive. Notice that for a given s we have to compute all parameters that depend on s before evaluating this function. Since it is not suitable to use the derivative of this function, a good choice for finding the minimum is to use Brent's algorithm (e.g. [FMM77]). To reduce the cost for finding a proper s, we do not minimize for every cone. Instead we use the following procedure: (1) Make some triangulation steps as described in x2.4. (2) Compute s for every cone C. (3) Continue with triangulation. When a cone is split by barycentric subdivision of the corresponding simplex, both new cones inherit s from the old simplex. Our computational experiences with various normal distributions show, that the costs for setup reduces enormously without raising the rejection constant too much. Using this procedure it might happen that s does not give a proper touching point (or HC is too big; see end of x2.4) after finishing all triangulation steps. Thus we have to check s for every cone and continue with triangulation in some cones if necessary. 3.2 Sampling The subroutine sample() consists of four parts: (S1) select a cone C, (S2) find a random variate proportional to the marginal density h g (27), (S3) generate a uniform random tuple U on the standard simplex (i.e. and compute tuple on the intersection Q(x) of the sweeping plane with cone C. (S3) and (S4) is done by subroutine simplex(). 4 Possible variants 4.1 Subset of R n as domain Our experiments have shown, that the basic algorithm works even for densities with support Since the hat h(x) has support the rejection constant might become very big. Subroutine 3 Construct a hat function Input: level of triangulation for finding s, level of minimal triangulation 1: for all tuples 2: Append new cone to list of cones with en as its spanning vectors. 3: repeat 4: for all cone C in list of cones do 5: call split() with C. Update list of cones. 7: until level of triangulation for finding s is reached 8: for all cone C in list of cones do 9: call find() with C. 10: repeat 11: for all cone C in list of cones do 12: call split() with C. 13: Update list of cones. 14: until minimum level of triangulation is reached 15: repeat 16: for all cone C in list of cones where s unknown do = (13) violated 17: call split() with C and list of cones. call find() with both new cones. 19: Update list of cones. 20: until no such cone was found 21: for all cone C in list of cones do 22: Compute all parameters of C. Total volume below hat 24: for all cone C in list of cones do Used for O(0)-search algorithm 27: return list of cones, H tot . Subroutine 4 cone and update list of cones Input: cone C, list of cones 1: Find lowest indices i; j of all vectors of C. 2: Find highest index m of all vectors (of triangulation). 3: 4: Append new cone C 0 to list and copy vectors and s of C into C 0 . 5: Replace t i by t m+1 in C and replace t j by t m+1 in C 0 . Replace det by 1 in C and C 0 . 7: return list of cones. Subroutine 5 find a proper touching point Input: cone C Bracketing a minimum 1: Search for a s 1 2 DA . return failed if not successful. 2: Search for s 0 , s 2 (Use property (21)). return failed if not successful. 3: Find s using Brent's algorithm (Use (29)). return failed if not successful. Subroutine 6 Generate a random tuple with density proportional to hat Input: H tot , list of cones 1: Generate a uniformly [0; H tot ] distributed random variate U . 2: Find C, such that H cum pred C . (C pred is the predecessor of C is the list of cones.) 3: Generate a gamma(n; fi) distributed random variate G. uniformly distributed point in Q(G) and return tuple 4: X / call simplex() with C and G. 5: return X. Subroutine 7 Generate a uniform distributed tuple on simplex Input: cone C, x (location of sweeping plane) uniformly distributed random variates U i in simplex 1: Generate iid uniform [0; 1] random variates U i , 2: Un / 1. 3: 4: for do 5: U i / U uniformly distributed point X in Q(x) x 7: return X. Pyramids. If the given domain D is a proper subset of R n (that is, we give constraints for suppf ), the acceptance probability can be increased when we restrict the domain of h accordingly to the domain D. (The domain is the set of points where the density f is defined; obviously suppf ' D. Notice that we have to provide the domain D for the algorithm but the support of f is not known.) Thus we replace (some) cones by pyramids. Notice that the base of such a pyramid must be perpendicular to the direction g. Hence we first have to choose a construction point p and then compute the height of the pyramid. The union of these pyramids (and of the remaining cones) must cover D. Whenever we get a random point not in the domain D we reject it. It is clear that continued triangulation decreases the volume between D and enclosing set. Polytopes. We only deal with the case where D is an arbitrary polytope which are given by a set of linear inequalities. Height of pyramid. The height is the maximum of hg; xi in C " D. Because of our restriction to polytopes this is a linear programming problem. Using the spanning vectors t as basis for the R n , it can be solved by means of the simplex algorithm in at most k pivot steps (for a simple polytope), where k is the number of constraints for D. Marginal density and volume below hat. The marginal distribution is a truncated gamma distribution with domain [0; u], where u is the height of the pyramid C. Instead of (28) and (29) we find for pyramids Z ux \Gamman \Gamma(n; fiu) (30) and where \Gamma(n; R x \Gammat dt is proportional to the incomplete gamma function and can be computed by means of formula (3.351) in [GR65]. Computing the height u(s) is rather expensive. So it is recommended to use instead of the exact function (31) for finding a touching point in pyramid C. Computational experiments with the standard normal distribution have shown, that the effect on the rejection constant is rather small (less than 5%). 4.2 Density not differentiable For the construction of the hat function we need a tangent plane for every x 2 D. Differentiability of the density is not really required. Thus it is sufficient to have a subroutine that returns the normal vector of a tangent hyperplane (which is However for densities f which are not differentiable the function in (29) might have a nasty behavior. However notice that f must be continuous in the interior of suppf , since log ffi f is concave. 4.3 Indicator Functions If is the indicator function of a convex set, then we can choose an arbitrary point in the convex set as the mode (as origin of our construction) and set t, the center of the cone C (see (4) in x2.2). Notice that the marginal density in (16) now reduces to h g . None of the parameters ff and fi depends on the choice of the touching point p. Of course we have to provide a compact domain for the density. Using indicator functions we can generate uniformly distributed random variates of arbitrary convex sets. 4.4 Mode not in Origin It is obvious that the method works, when the mode m is an arbitrary point in R. If the mode is unknown we can use common numerical methods for finding the maximum of f , since T (f(x)) is concave (see e.g. [Rao84]). Notice that the exact location of the mode is not really required. The algorithm even works if the center for the construction of the cones is not close to the mode. Then we just get a hat with a worse rejection constant. 4.5 Add mode as construction point Since we have only one construction point in each cone, the rejection constant is bounded from below. Thus only a few steps to triangulate the S sense. To get a better hat function we can use the mode m of f as an additional construction point. The hat function is then the minimum of f(m) and the original hat. The cone is split into two parts by a hyperplane F (b) with different marginal densities, where b is given by T f(m). Its marginal density is then given by Notice that we use the same direction g for the sweep plane in both parts. We have to compute the volume below the hat for both parts which are given by a x 4.6 More construction points per cone A way to improve the hat function is to use more than one (or two) construction points. But this method has some disadvantages and it is not recommended to use it. overlapping region Figure 3: Two construction points in a cone The cones are divided in several pieces of a pyramid (see figure 3). The lower and upper base of these pieces must be perpendicular to the corresponding direction g. These vectors g are determined by the gradients of the transformed density at the construction points in this pieces. Thus these g (may) differ and hence these pieces must overlap. This increases the rejection constant. Moreover it is not quite clear how to find such pieces. For the univariate case appropriate methods exist (e.g. [DH94]). But in the multivariate case these are not suitable. Also adaptive rejection sampling (introduced in [GW92]) as used in [H-or95b, LH98] is not a really good choice. The reason is quite simple. The cones are fixed and the construction points always are settled in the center of these cones. Thus using adaptive rejection sampling we select the new construction points due to a distribution which is given by the marginal density of this marginal density is not zero at the existing construction points. 4.7 Squeezes We can make a very simple kind of squeezes: Let x Compute the minima of the transformed density at Q(x i ) for all i. Since ~ f is concave these minima are at the vertices of these simplices. The squeeze s i (x) for x denotes the minimum of ~ f(x) in Q(x i ). The setup of these squeezes is rather expensive and only useful, if many random points of the same distribution must be generated. -concave densities A family T c of transformations that fulfil conditions (T1)-(T4) is introduced in [H-or95a]. Let c - 0. Then we set c (x) It can easily be verified, that condition (T4) (i.e. volume below hat is bounded) holds if and only if \Gamma 1 To ensure the negativity of the transformed hat we always have to choose the mode m as construction point (see x4.5). In [H-or95a] it was shown that if a density f is T c -concave then it is T c1 - concave for all c 1 - c. The case log(x) is already described in x2.6. For the case c ! 0 the marginal density function (16) is now given by c for x ? b where b is given by (fi To our knowledge no special generator for this distribution is known. (The part for x ? b looks like a beta-prime distribution (see [JKB95]), but ff; fi ? 0. By assumption (fi x \Gamman. Hence it can easily be seen that the marginal density is T c -concave. Therefore we can use the universal generator ([H-or95a]). Computational Experience 5.1 A C-implementation. A test version of the algorithm was written in C and is available via anonymous ftp [Ley98]. It should handle the following densities ffl f is log-concave but not constant on its support. ffl Domain D is either R n or an arbitrary rectangle [a ffl The mode m is arbitrary. But if D 6= suppf then m must be an interior point of suppf not "too close" to the boundary of suppf . We used two lists for storing the spanning vectors and the cones (with pointers to the list of vectors). For the setup we have to store the edges (i; computing the new vertices. This is done temporarily in a hash table, where the first index i is used as the hash index. The setup step is modified in the case of a rectangular domain. If the mode is near to the boundary of D we use the nearest point on the boundary (if possible a vertex) for the center to construct the cones. If this point is on the boundary we easily can eliminate all those initial cones, that does not intersect D. If this point is a vertex of the rectangle there remains only one initial cone. For finding the mode of f we used a pattern search method by Hooke and Jeeves [HJ61, Rao84] as implemented in [Kau63, BP66], since it could deal with both unbounded and bounded support of f without giving explicit con- straints. For finding the minimum of (29) we use Brent's algorithm as described in ([FMM77]). The implementation contains some parameters to adjust these routines to a particular density f . For finding a cone C in subroutine sample() we used a O(0)-algorithm with a search table. (Binary search is slower.) For generating the gamma distributed random number G we used the algorithm in [AD82] for the case of unbounded domain. When D is a rectangle, we used transformed density rejection ([H-or95a]) to generate from the truncated gamma distributions. Here it is only necessary to generate a optimal hat function for the truncated gamma distribution with shape parameters n and 1 with domain (0; umax ), where umax is the maximal value of height \Delta fi for all cones. The optimal touching points for this gamma distribution are computed by means of the algorithm [DH94]. The code was written for testing various variants of the algorithm and is not optimized for speed. Thus the data shown in the tables below give just an idea of the performance of the algorithm. We have tested the algorithm with various multivariate log-concave distributions in some dimensions. All our tests have been done on a PC with a P90 processor running Linux and the GNU C compiler. 5.2 Basic version: unbounded domain, mode in origin Random points with density proportional to hat function. The time for the generation of random points below the hat has shown to be almost linear in dimension n. Table 1 shows the average time for the generation of a single point. For comparison we give the time for generation of n normal distributed points using the Box/Muller method [BM58] (which gives a standard multi-normal distributed point with density proportional to For computing the hat function we only used initial cones for the standard multinormal density. hat function 14.6 17.1 21.3 24.9 30.2 34.6 41.5 45.7 55.6 multinormal 7.2 10.8 14.4 18.0 21.6 25.2 28.8 32.4 36.0 Table 1: average time for the generation of one random point (in -s) Random points for the given density. The real time needed for the generation of a random point for a given log-concave density depends on the rejection constant and the costs for computing the density. Table 2 shows the acceptance probabilities and the times needed for the generation of standard multinormal distributed points. Notice that these data do not include the time for setting up the hat function. Setup. When find() is called after triangulation has been done, the time needed for the computation of the hat function depends linearly on the number number of cones time (-s) 23.7 27.9 38.2 49.8 73.8 99.3 142 262 575 acceptance Table 2: acceptance probability and average time for the generation of standard multinormal distributed points of cones. (Thus find() is the most expensive part of the setup().) Table 3 shows the situation for the multinormal distribution with density proportional to It demonstrates the effects of continuing barycentric subdivision of the "oldest" edge (see x2.4) on the number of cones, the acceptance probability, the costs for generating a random point proportional to the hat function (i.e. without rejection) and proportional to the given density. Furthermore it shows the total time (in ms) for the setup (i.e. for computing the parameters of the hat function) (in ms) and the time for each cone (in -s) (The increase for large n in the time needed for generating points below the hat is due to effects of memory access time.) subdivisions cones acceptance hat (-s) 24.8 24.8 24.9 25.0 25.2 25.3 25.7 26.1 27.1 27.6 28.4 density (-s) 94.2 73.0 59.9 52.1 45.6 42.9 40.1 39.3 39.5 39.6 40.4 setup/cone (-s) 700 706 738 720 713 714 727 756 762 763 764 Table 3: time for computing the hat function for multinormal distribution If we do not run find() for every cone of the triangulation but use the method described in x3.1 we can reduce the costs for the construction of the hat function. Table 4 gives an idea of this reduction for the multinormal distribution with proportional to It shows the time for constructing the hat function subject to the number of cones for which find() is called. Due to Table 4 the acceptance probability is not very bad, if we run find() only for the initial cones. But this is not true in general. It might become extremely poor if the level sets of the density are very "skinny". Table 5 demonstrates the effect on the density proportional to 2. At last table 6 demonstrates that the increase in the time for constructing the hat function for increasing dimension n is mainly due to the increase of find() in subdivision cones acceptance setup (ms) 66.2 76.8 100.2 141.6 224.7 393.6 744 setup/cone (total) (-s) Table 4: time for computing the hat function for multinormal distribution with "inherited" construction points find() in subdivision 1 2 3 4 acceptance Table 5: acceptance probability for multinormal distribution with "inherited" construction points number of cones. Notice that we start with 2 n cones. Furthermore we have to make consecutive subdivisions to shorten every edge of a simplex that defines an initial cone. Thus the number of cones increases exponentially. acceptance setup/cone (-s) 540 616 714 811 927 1011 1170 1250 1421 Table time for computing the hat function for multinormal distribution subdivisions of the initial cones) If the covariance matrix of the multinormal distribution is not a diagonal matrix and the ratio of the highest and lowest eigenvalue is large, then we cannot use initial cones only and we have to make several subdivisions of the cones. Because of the above considerations the necessary number of cones explodes with increasing n. Thus in this case this method cannot be used for large n. (Suppose we have to shorten every edge of each simplex, then we have cones if but we need 2 Tests. We ran a - 2 -test with the density proportional to exp(\Gamma to validate the implementation. For all other densities we compared the observed rate of acceptance to the expected acceptance probability. Comparison with algorithm [LH98]. The code for algorithm [LH98] is much longer (and thus contains more bugs). The setup is much slower and it needs 11 750 -s to generate on mulitnormal distributed random point in dimension 4 (versus 38 -s in table 2 for tdrmv()). 5.3 Rectangular domain Normal densities restricted to an arbitrary rectangle have a similar performance as the corresponding unrestricted densities. except of the acceptance probability which is worse since the domain of the hat h is a superset of the domain of density f . 5.4 Quality The quality of non-uniform random number generators using transformation techniques is an open problem even for the univariate case (see e.g. [H-or94] for a first approach). It depends on the underlying uniform random number generators. The situation is more serious in the multivariate case. Notice that this new algorithm requires more than n+2 uniform random numbers for every random point. We cannot give an answer to this problem here, but it should be clear that e.g. RANDU (formerly part of IBM's Scientific Subroutine Package, and now famous for its devasting defect in three dimensions: its consecuting points just fifteen parallel planes; see e.g. [LW97]) may result in a generator of poor quality. 5.5 Some Examples We have tested our algorithm in dimensions proportional to where a i ? 0. The domain was R n and some rectangles. We also used densities proportional to f i (U x U is an orthonormal transformation and b a vector, to test distributions with non-diagonal correlation matrix and arbitrary mode. The algorithm works well for densities f 3 , f 4 and f 5 both with and D being a rectangle enclosing the support of f i . Although some of these densities are not C 1 , the find() routine works. Problems arise if the level sets of the density have "corners", i.e. the g is unstable when we vary the touching a little bit. Then there are somes (that contains these "corners") with huge volumeHC and further triangulation is necessary. If dimension is high (n & 5) too many cones might be necessary. The optimization algorithm for finding the mode fails if we use a starting point outside the support of f 5 . The code has some parameters for adjusting the algorithm to the given density. For example, it requires some testing to get the optimal number of cones and the optimal level of subdivisions for calling find(). 5.6 R'esum'e The presented algorithm is a suitable method for sampling from log-concave (and T -concave) distributions. The algorithm works well for all tested log-concave densities if dimension is low (n . 5) or if correlation is not too high. Restrictions of these densities to compact polytopes are possible. The setup time is small for small dimension but increases exponentially in n. The speed for generating random points is quite fast even for n - 6. Due to the large amount of cones for high dimension the program requires a lot of computer memory (typically 2-10 MB). Although the developed algorithm is not a real black box method it is easily adjustable for a large class of log-concave densities. Examples for which the algorithm works are the multivariate normal distribution and the multi-variate student distribution (with transformation T arbitrary mean vector and variance matrix conditioned to an arbitrary compact polytope. However for higher dimensions the ratio of highest and lowest eigenvalue of the covariance matrix should not be "too big". Acknowledgments The author wishes to note his appreciation for help rendered by J-org Lenneis. He has given lots of hints for the implementation of the algorithm. The author also thanks Gerhard Derflinger and Wolfgang H-ormann for helpful conversations and their interest in his work. --R Generating gamma variates by a modified rejection technique. Box and M. Remark on algorithm 178. Principles of Random Variate Generation. Random variate generation for multivariate densities. The optimal selection of hat functions for rejection algorithms. Random variable generation using concavity properties of transformed densities. Computer methods for mathematical computations. Table of Integrals. Convex Poytopes. Adaptive rejection sampling for gibbs sam- pling Universal generators for correlation induction. "Direct search" The quality of non-uniform random numbers A rejection technique for sampling from T A universal generator for bivariate log-concave distri- butions Continuous Univariate Distributions Multivariate Statistical Simulation. Polytope Volume A sweep plane algorithm for generating random tuples. Inversive and linear congruential pseudorandom number generators in empirical tests. Theory and Applications. On computer generation of random vectors by transformations of uniformly distributed vectors. The Computation of Fixed Points and Applications Improving the convergence of fixed-point algorithms Efficient generation of random variates via the ratio-of-uniforms method Lectures on Polytopes --TR Multivariate statistical simulation On computer generation of random vectors by transformation of uniformly distributed vectors A rejection technique for sampling from <italic>T</italic>-concave distributions Inversive and linear congruential pseudorandom number generators in empirical tests Random variate generation for multivariate unimodal densities A sweep-plane algorithm for generating random tuples in simple polytopes `` Direct Search'''' Solution of Numerical and Statistical Problems Generating gamma variates by a modified rejection technique Remark on algorithm 178 [E4] direct search Algorithm 178: direct search Computer Methods for Mathematical Computations --CTR G. Leobacher , F. Pillichshammer, A method for approximate inversion of the hyperbolic CDF, Computing, v.69 n.4, p.291-303, December 2002 Wolfgang Hrmann, Algorithm 802: an automatic generator for bivariate log-concave distributions, ACM Transactions on Mathematical Software (TOMS), v.26 n.1, p.201-219, March 2000 W. Hrmann , J. Leydold, Random-number and random-variate generation: automatic random variate generation for simulation input, Proceedings of the 32nd conference on Winter simulation, December 10-13, 2000, Orlando, Florida Seyed Taghi Akhavan Niaki , Babak Abbasi, Norta and neural networks based method to generate random vectors with arbitrary marginal distributions and correlation matrix, Proceedings of the 17th IASTED international conference on Modelling and simulation, p.234-239, May 24-26, 2006, Montreal, Canada sampling with the ratio-of-uniforms method, ACM Transactions on Mathematical Software (TOMS), v.26 n.1, p.78-98, March 2000

rejection method;multivariate log-concave distributions

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Mixed-Mode BIST Using Embedded Processors.

In complex systems, embedded processors may be used to run software routines for test pattern generation and response evaluation. For system components which are not completely random pattern testable, the test programs have to generate deterministic patterns after random testing. Usually the random test part of the program requires long run times whereas the part for deterministic testing has high memory requirements.In this paper it is shown that an appropriate selection of the random pattern test method can significantly reduce the memory requirements of the deterministic part. A new, highly efficient scheme for software-based random pattern testing is proposed, and it is shown how to extend the scheme for deterministic test pattern generation. The entire test scheme may also be used for implementing a scan based BIST in hardware.

Introduction Integrating complex systems into single chips or implementing them as multi-chip modules (MCMs) has become a widespread approach. A variety of embedded processors and other embedded coreware can be found on the market, which allows to appropriately split the system functionality into both hardware and software modules. With this development, however, system testing has become an enormous challenge: the complexity and the restricted accessibility of hardware components require sophisticated test strategies. Built-in self-test combined with the IEEE 1149 standards can help to tackle the problem at low costs [10]. For conventional ASIC testing, a number of powerful BIST techniques have been developed in the past [1 - 3, 5, example, it has been shown that combining random and efficiently encoded deterministic patterns can provide complete fault coverage while This work has been supported by the DFG grant "Test und Synthese schneller eingebetteter Systeme" (Wu 245/1-2). keeping the costs for extra BIST hardware and the storage requirements low [13, 14, 32]. In the case of embedded systems such a high quality test is possible without any extra hardware by just using the embedded processor to generate the tests for all other components. Usually, this kind of functional testing requires large test programs, and a memory space not always available on the system. In this paper it will be shown how small test programs can be synthesized such that a complete coverage of all non-redundant stuck-at faults in the combinational parts of the system is obtained. The costs for extra BIST hardware in conventional systems testing are reduced to the costs for some hundred bytes of system memory to store the test routines. The proposed BIST approach can efficiently exploit design-for-testability structures of the subcomponents. As shown in Figure 1 during serial BIST the embedded processor executes a program which generates test patterns and shifts them into the scan regis- ter(s) of the component(s) to be tested. Even more effi- ciently, the presented approach may be used to generate test data for input registers of pipelined or combinational subsystems. embedded processor scan- input scan- output scan- input scan- output test data (random & deterministic patterns) Figure 1: Serial BIST approach. The structure of the test program can be kept very sim- ple, if only random patterns have to be generated, since then some elementary processor instructions can be used [12, 21, 25, 28]. Even a linear feedback shift register (LFSR) can be emulated very efficiently: Figure 2 shows as an example a modular LFSR and the corresponding program (for simplicity the C-code is shown) to generate a fixed number of state transitions. void transition (int m, int n, unsigned int polynomial, unsigned int *state) transitions of modular LFSR of degree n */ { int for (i=0; i<m; i++) { if (*state >> n-1) { *state <<= *state ^= polynomial; else *state <<= Figure 2: Modular LFSR and corresponding program for generating state transitions. But usually not all the subcomponents of a system will be random pattern testable, and for the remaining faults deterministic test patterns have to be applied. For this pur- pose, compact test sets may be generated as described in [16, 18, 22, 27] and reproduced by the test program, or a hardware-based deterministic BIST scheme is emulated by the test software [13 - 15, 32]. This kind of mixed-mode testing may interleave deterministic and random testing or perform it successively. In each case, the storage requirements for the deterministic part of the test program are directly related to the number of undetected faults after random pattern generation. There is a great trade-off between the run-time for random test and the memory requirements of the mixed-mode program. Assume a small improvement of the random test method which leads to an increase of the fault coverage from 99.2% to 99.6%. This reduces the number of undetected faults and the storage requirements by the factor 1/2. Overall, the efficiency of a mixed-mode test scheme can be improved to a much higher degree by modifying its random part rather than its deterministic part. In this paper a highly efficient software-based random BIST scheme is presented which is also used for generating deterministic patterns. The rest of the paper is organized as follows: In the next section, different random pattern test schemes to be emulated by software are evaluated, and in section 3 the extension to deterministic testing is described. Subsequently, in section 4, a procedure for optimizing the overall BIST scheme is presented, and section 5 describes the procedure for generating the mixed-mode test program. Finally, section 6 gives some experimental results based on the INTEL 80960CA processor as an example Emulated Random Pattern Test Test routines exploiting the arithmetic functions of a processor can produce patterns with properties which are sufficient for testing random pattern testable circuits [12, 25], even if they do not completely satisfy all the conditions for randomness as stated in [11], e.g. However, for other circuits, in particular for circuits considered as random pattern resistant, arithmetic patterns may not perform as well. Linear feedback shift registers (LFSRs) corresponding to primitive feedback polynomials and cellular automata are generally considered as stimuli generators with good properties for random testing [9, 17, 20]. But the generated sequences still show some linear dependen- cies, such that different primitive polynomials perform differently on the same circuit. In some cases, the linear dependencies may support fault detection, for other circuits they perform poorly. In the following, the fault coverage obtained by several LFSR-based pattern generation schemes will be discussed with some experimental data. 2.1 Feedback Polynomial In contrast to hardware-based BIST, in a software-based approach the number and the positions of the feedback taps of the LFSR have no impact on the costs of the BIST im- plementation. Thus, for a given length the achievable fault coverage can be optimized without cost constraints. Assuming a test per scan scheme as shown in Figure 3 the sensitivity of the fault coverage to the selected feed-back polynomial has been studied by a series of experiments for the combinational parts of the ISCAS85 and ISCAS89 benchmark circuits [4, 6]. LFSR feedback scan path CUT Figure 3: Scan-based BIST. Circuit PI F Degree LFSR1 LFSR2 LFSR3 LFSR4 LFSR5 LFSR6 Average Table 1: Absolute and normalized (w. r. t. worst LFSR) percentage of undetected non-redundant faults after 10,000 patterns. Fault simulation of 10,000 random patterns was performed for each circuit using several different feedback polynomials, all of the same degree. Some typical results are shown in Table 1. The first four columns contain the circuit name, the number of inputs, the number of non-redundant faults, and the selected degree of the feedback polynomial. 1 The remaining columns show the characteristics for six different LFSRs. The first entry reports the percentage of undetected non-redundant faults, and the second entry normalizes this number to the corresponding number for the worst LFSR (in %). The worst and best performing LFSR are printed in bold, respectively. The last column gives the average over all of the LFSRs. It can be observed that there is a big variance in the performance of different LFSRs of the same degree. For s641, e.g., the best LFSR reduces the number of undetected faults down to 27% of the faults left undetected by the worst polynomial. 2.2 Multiple-Polynomial LFSRs One explanation for the considerable differences in fault coverage observed in section 2.1 is given by the fact, that linear dependencies of scan positions may prevent certain necessary bit combinations in the scan patterns independent of the initial state of the LFSR [2]. For different LFSRs the distribution of linear dependencies in the scan chain is different and, depending on the structure of the circuit, may have a different impact on the fault coverage. As shown in Figure 4 the impact of linear dependencies can be reduced if several polynomials are used. In this small example the LFSR can operate according to two dif- 1 The degrees of the polynomials have been selected, such that they were compatible with the requirements for the deterministic test described in section 3. ferent primitive feedback polynomials P and are selected by the input of the multiplexer. For any given initial state LFSR produces a scan pattern (a 0 , . , a 7 ), such that, depending on the selected polynomial, the shown equations for hold for its components. st-0 st-0 a 0 a 1 a 2 a 3 a 4 a 5 a 6 a 7 U I Figure 4: Scan-based BIST with multiple-polynomial LFSR. For polynomial P 0 there is a linear relation a 3 which prevents the combination (1, 1, 1) at the inputs of the AND-gate. This implies that the polynomial P 0 (X) can never produce a test for the stuck-at-0 fault at node o 2 . In contrast to that, for polynomial the same input positions are linearly independent and produce all possible nonzero bit combinations and thus a test for the considered fault. Similarly, the stuck-at-0 fault at node o 1 cannot be tested using polynomial polynomial P 0 (X) can provide a test. Using both polynomials, each for a certain number of patterns, increases the chance of detecting both faults. Such a multiple-polynomial LFSR can be implemented efficiently in hardware by trying to share parts of the feed-back for several polynomials. A software emulation is also very simple, since the basic procedure to simulate an LFSR has to be modified only slightly. To control the selection of feedback polynomials several schemes are possible. The first is shown in Figure 5 assuming N random patterns to be generated by p different polynomials denotes the LFSR operation corresponding to feedback polynomial P i . initialize (LFSR); for generate -N/p- patterns by LFSR(P i ); Figure 5: Successive multiple-polynomial scheme (SUC). The polynomials are applied successively to generate contiguous subsequences of -N/p- random patterns, the scheme will therefore be referred to as scheme SUC. For one polynomial the scheme degenerates to the conventional single polynomial scheme. The possibility to switch between different distributions of linear dependencies is paid by the disadvantage that some patterns may occur repeatedly up to p times. Hence, an overall increase of the fault coverage cannot be expected, but experiments have shown that there is indeed an improvement for some circuits. Table 2 lists the results for the same set of circuits as studied in the previous section. Circuit Degree Table 2: Absolute and normalized (w. r. t. worst and best single LFSR) percentage of undetected non-redundant faults for scheme SUC after 10,000 patterns. For each circuit 10,000 patterns were simulated using p polynomials. For each experiment the percentage of undetected non-redundant faults is reported (1st line), as well as the corresponding normalized numbers for the worst (2nd line) and for the best single polynomial (3rd line) of the same degree (in %). Applying the successive scheme for example to the circuit c2670 with reduces the number of undetected faults down to 69.58% compared with the worst single polynomial. Even more important is that the scheme also outperforms the best single polynomial and the number of remaining target faults for ATPG is less than 75%, i.e. 25% percent of the faults left by the best single polynomial are additionally covered by this scheme. The randomness of the sequence can be further in- creased, if the polynomials are not used successively, but selected randomly for each test pattern. This random selection can be implemented by a second LFSR as shown in Figure 6 and will be referred to as scheme RND. U feedback CUT scan chain Figure Hardware scheme for the random selection of feed-back polynomials (RND). The selection between p different feedback polynomials for LFSR1 is controlled by -log 2 p- bits of the state register of LFSR2. For a software implementation of the structure of Figure 6, two additional registers are required for storing the feedback polynomial and the state of LFSR2. and LFSR2 can be emulated by the same proce- dure, and the complete routine to generate a sequence of N random patterns is shown in Figure 7. initialize (LFSR1); initialize (LFSR2); for { select P based on state of LFSR2; generate 1 pattern by LFSR1(P); perform 1 state transition of LFSR2; Figure 7: Software routine for the random pattern generation scheme of Figure 6 (RND). Table 3 shows the percentage of undetected non-redundant faults and the corresponding normalized numbers obtained by the scheme RND for Circuit Degree Table 3: Absolute and normalized (w. r. t. worst and best single LFSR) percentage of undetected non-redundant faults for scheme RND after 10,000 patterns. For the randomly selected polynomials, there is a higher chance of pattern repetitions, but randomly switching between different distributions of linear dependencies may improve the quality of the patterns. For some cir- cuits, this results in an improvement of fault coverage, so that the set of faults which remain for deterministic testing is further reduced. 2 . 3 Multiple-Polynomial , Multiple-Seed Another way of improving the efficiency of a random test is repeatedly storing a new seed during pattern generation as investigated for instance in [23]. This technique can be combined with the use of multiple polynomials as shown in Figure 8. U feedback CUT scan chain Figure 8: Multiple-polynomial, multiple-seed LFSR. As for the scheme RND, -log 2 p- bits of the state register of LFSR2 are used to drive the selection between p different feedback polynomials of degree k for LFSR1. The remaining k bits provide the seed for LFSR1. In the sequel this scheme will be referred to as the scheme RND 2 . The structure of the corresponding test program is shown in Figure 9. initialize (LFSR2); for { select seed S and polynomial P based on state of LFSR2; initialize LFSR1 with generate 1 pattern by LFSR1(P); perform 1 state transition of LFSR2; Figure 9: Test program for the multiple-polynomial, multi- ple-seed LFSR (RND 2 ). Again, in this scheme patterns may occur repeatedly, but in addition to the advantage of randomly changing the distribution of linear dependencies this scheme is also able to generate the all zero-vector which is often needed for complete fault coverage. Table 4 gives the results for (percentage of undetected non-redundant faults and the corresponding normalized numbers as in Tables 2 and 3). Circuit Degree Table 4: Absolute and normalized (w. r. t. worst and best single LFSR) percentage of undetected non-redundant faults for scheme RND 2 after 10,000 patterns. As expected, not for all circuits the fault coverage increases, but there are circuits where this technique leads to significant improvements. For circuits s838.1 and s9234 the best results are obtained compared with all the experiments before. 3 Software-Based Deterministic BIST The structure of the multiple-polynomial, multiple-seed random BIST scheme of Figure 8 is very similar to the deterministic BIST scheme based on reseeding of multiple- polynomial LFSRs proposed in [13, 14], see Figure 10. U CUT scan chain (m bits) id seed Figure 10: Deterministic BIST scheme based on a multiple- polynomial LFSR by [14]. A deterministic pattern is encoded as a polynomial identifier and a seed for the respective polynomial. During test mode the pattern can be reproduced by emulating the LFSR corresponding to the polynomial identifier, loading the seed into the LFSR and performing m autonomous transitions of the LFSR. After the m-th transition the scan chain contains the desired pattern which is then applied to the CUT. To calculate the encoding systems of linear equations have to be solved. For a fixed feedback polynomial of degree k the LFSR produces an output sequence (a i ) i-0 satisfying the feedback equation a k. The LFSR-sequence is compatible with a desired test pattern t specified bits a holds. Recursively applying the feedback equation provides a system of linear equations in the seed variables a 0 , ., a k-1 . If no solution can be found for the given polynomial, the next available polynomial is tried, and in [14] it has been shown that already for 16 polynomials there is a very high probability of success that a deterministic pattern with s specified bits can be encoded into an s-bit seed. Hence, if p different polynomials are available and the polynomial identifier is implemented as a "next bit", the seed and the next bits for a deterministic test set number of specified bits s max require bits of storage. Minimizing S(T) requires both minimizing the maximum number of care bits s max and the number of patterns N. In [15] an ATPG-algorithm was presented which generates test patterns where the number of specified bits s max is mini- mized. In a mixed-mode BIST approach the number N of patterns is highly correlated to the number of faults left undetected after random testing. 4 Synthesizing the BIST Scheme Since the efficiency of a mixed-mode BIST scheme strongly depends on the number of hard faults to be covered by deterministic patterns, a major concern in synthesizing the BIST scheme is optimizing the random test. The experimental data of section 2 show that significant variances in the fault efficiency achieved by different LFSR schemes exist, and that there is no universal scheme or polynomial working for all of the circuits. In the sequel, a procedure is presented for determining an optimized LFSR scheme. The selection of the LFSRs is guided, such that the fault efficiency is maximized while satisfying the requirements for an efficient encoding of deterministic patterns for the random pattern resistant faults. Assuming a table of primitive polynomials available the proposed procedure consists of 4 steps: Perform ATPG to eliminate the redundant faults and to estimate the maximum number of specified bits, s max , to be expected in the test cubes for the hard faults. Select M polynomials of degree s max randomly, and perform fault simulation with the corresponding shift register sequences. Rank the polynomials according to the fault coverage achieved. Select the P best polynomials and store the highest fault coverage and the corresponding LFSR as BEST_SCHEME. Using polynomials, simulate the schemes SUC, RND, and RND 2 . Update BEST_SCHEME to the best solution obtained so far. The number M is mainly determined by a limit of the computing time to be spent. The number P is also restricted by the computing time available, but in addition to that each LFSR requires two registers of the processor for pattern generation. So, the register file of the target processor puts a limit on P, too. Table 5 shows the results achieved by this procedure for the same set of circuits as studied in section 2. For the same degrees as used in section 2 sequences of 10,000 random patterns were applied. Scheme best UF worst Table 5: Best schemes and relation to best and worst single polynomial solution. The second and third column show the best scheme and the corresponding number of polynomials p, column 4 provides the fault efficiency FE (percentage of detected non-redundant faults). The percentage of faults left undetected by the best scheme is reported in column UF. UF best normalizes this solution to the number obtained by the best single polynomial, UF worst refers to the worst single polynomial. Table 5 indicates that the search for an appropriate random test scheme can reduce the number of remaining faults significantly. The procedure needs M runs of fault simulations, but may decrease the storage amount needed for deterministic patterns considerably. These savings in memory for the mixed-mode test program are particularly important, if the test program has to be stored in a ROM for start-up and maintenance test. Generating Mixed-Mode Test Programs Test programs implementing the random test schemes and the reseeding scheme for deterministic patterns were generated for the INTEL 80960CA as a target processor. Its large register set made a very compact coding possible. Since the part of the test program which generates the deterministic patterns is a superset of instructions required for implementing any of the random schemes, only the example for the most complex random scheme is shown. The mixed-mode test program of Figure 11 generates random test patterns by multiple-polynomial, multiple- seed LFSR emulation, and switches to the reseeding scheme afterwards. The program of Figure 11 requires 27 words in memory but assumes that all LFSRs fit into registers. This steps1 equ . ; number of steps for lfsr1 steps2 equ . ; number of steps for lfsr2 steps_det equ . ; number of steps for deterministic test len1 equ . ; position of msb of lfsr1 len2 equ . ; position of msb of lfsr2 testport equ . ; address of testport no_poly_bits equ . ; number of bits for polynomial choice mask equ . , define mask start dq startvector ; define startvector for lfsr2 poly dq polynomials ; define polynomials for lfsr1 seeds dq seedvectors ; define seeds for det. test seed_offset equ seeds - start ; define offset for seed table begin: lda testport, r10 ; load address of testport lda steps_det, r11 ; load loopcounter for lfsr1 in det. mode lda steps1, r12 ; load loop counter for lfsr1 lda start, r14 ; load startvector address for lfsr1 ld (r14), r6 ; load startvector for lfsr2 ld 4(r14), r7 ; load polynomial for lfsr2 l0: mov r6, r4 ; initialize lfsr1 with contents of lfsr2 and mask, r4, r15 ; compute poly-id ld 8(r14)[r15*4], r5 ; polynomial for lfsr1 lda no_poly_bits, r15 ; load number of bits for poly-id l1: shro no_poly_bits, r4, r4 ; shift poly-bits lda steps2, r13 ; load loop counter for lfsr1 l2: st r4, (r10) ; write testpattern to testport mov r4, r8 shlo 1, r8, r4 ; shift left bbc len2, r8, l3 ; branch if msb of lfsr2 equal zero xor r4, r5, r4 ; xor decrement loop counter cmpibne r13, 0, l2 ; branch not equal zero mov r6, r8 shlo 1, r8, r6 ; shift left bbc len1, r8, l4 ; branch if msb of lfsr1 equal zero xor r6, r7, r6 ; xor l4: subi r12, 1, r12 ; decrement loop counter cmpibg r12, r11, l0 ; branch if r12 > steps_det ld seed_offset(r14)[r12*4],r6 ; load seed cmpibne r12,0,l0 Figure 11: Mixed-mode BIST program. is always possible for random pattern generation, but encoding deterministic patterns may lead to LFSR lengths exceeding bits. In this case, the program of Figure 11 has to be modified in a straightforward way, and requires more memory. Table 6 gives the relation between memory requirements and LFSR lengths. LFSR length Memory requirements (words) Table length and memory requirements for the mixed-mode test program. In addition to the program size, memory has to be reserved for storing the polynomials and the seeds in order to decode the deterministic patterns. The experimental results of the next section show that these data form by far the major part of the memory requirements. 6 Experimental Results The described strategy for generating mixed-mode test programs was applied to all the benchmark circuits for M e. for each circuit M 28 runs of fault simulation were performed to determine the best random scheme. Tables 7 and 8 show the results. Circuit PI Degree Best Scheme p Table 7: Circuit characteristics and best random scheme. The selected random schemes and their characteristic data are reported in Table 7. Columns 2 and 3 list the number of primary inputs PI and the degree of the poly- nomials. The best random scheme and the number of polynomials are reported in the subsequent columns. Table 8 shows the detailed results. The number of non-redundant faults for each circuit is given in column 2. The efficiency of the random scheme is characterized again by the fault efficiency FE, the percentage of undetected non-redundant faults UF and the normalized numbers for UF with respect to the best (UF best ) and the average (UF average ) single polynomial solution in columns 3 through 6. Circuit F FE UF UF best UF average s838.1 931 76.48 23.52 71.1 65.75 Table 8: Fault efficiency and percentage of undetected non-redundant faults for the best random schemes after 10,000 patterns. The reduction of the remaining faults obtained by the best random test scheme is significant. For instance, the circuit c7552 is known to be very random pattern resis- tant, and a single polynomial solution in the average leads to a fault efficiency of 95.79% leaving 4.21% of the faults for deterministic encoding. For the same circuit, the RND 2 scheme achieves a fault efficiency of 98.87%, and only 1.13% or, absolutely, 84 faults are left. This corresponds to a reduction of the remaining faults down to 27%. For circuits s820 and s1423 a careful selection of the random scheme even makes the deterministic test super- fluous. Finally, it should be noted that for the larger cir cuits already a small relative reduction means a considerable number of faults which are additionally covered by the random test and need not be considered during the deterministic test. For example for circuit s38417 a reduction down to 85.75% and 92.26%, respectively, means that additional 313 and 158, respectively, faults are eliminated during random test. Table 9 shows the resulting number of test patterns required for the random pattern resistant faults and the amount of test date storage (in bits) for the best random scheme compared to a random test using an average single polynomial. This includes the storage needed for the poly- nomials, the initial LFSR states for the random test and the encoded deterministic test set. Since the goal of this work was to determine the impact of the random test on the test data storage, a standard ATPG tool was selected to perform the experiments [24]. For all circuits the fault efficiency is 100% after the deterministic test. Circuit Deterministic patterns Test data storage (bits) scheme Average polynomial scheme Average polynomial s420.1 22 34 503 776 s1238 7 21 198 431 s5378 22 31 759 883 Table 9: Number of deterministic patterns and storage requirements for the complete test data (in bits). The results show that an optimized random test in fact considerably reduces the number of deterministic patterns and the overall test data storage. This is particularly true for the circuits known as random pattern resistant. E.g. for circuit c7552 the number of deterministic patterns is reduced from 92 to 51 and the reduction in test data storage is about 5K. For circuit s38417 the best scheme eliminates 137 deterministic patterns, which leads to a reduction in test data storage of more than 14K. As shown in Table already with standard ATPG the proposed technique requires less test data storage than an approach based on storing a compact test set (cf. [16, 18, 22, 27]). Circuit Deterministic patterns Test data storage (bits) scheme Compact Test scheme Compact Test s420.1 22 43 503 1505 s5378 22 104 759 22256 Table 10: Amount of test data storage for the proposed approach and for storing a compact test set. It can be expected, that the test data storage for the presented approach could be reduced even further, if an ATPG tool specially tailored for the encoding scheme were used as described in [15]. 7 Conclusion A scheme for generating mixed-mode test programs for embedded processors has been presented. The test program uses both new, highly efficient random test schemes and a new software-based encoding of deterministic patterns. It has been shown that the careful selection of primitive polynomials for LFSR-based random pattern generation has a strong impact on the number of undetected faults, and a multiple-polynomial random pattern scheme provides significantly better results in many cases. The quality 0of the random scheme has the main impact on the overall size of a mixed-mode test program. As an example, for the processor INTEL 80960CA test programs were generated, and for all the benchmark circuits a complete coverage of all non-redundant faults was obtained. --R Test Embedding in a Built-in Self-Test Environment Exhaustive Generation of Bit Patterns with Applications to VLSI Self-Testing A Neutral Netlist of 10 Combinational Benchmark Designs and a Special Translator in Combinational Profiles of Sequential Benchmark Circuits A New Pattern Biasing Technique for BIST BIST Hardware Generator for Mixed Test Scheme Multichip Module Self-Test Provides Means to Test at Speed Shift Register Sequences Test Generation Based On Arithmetic Operations Generation of Vector Patterns Through Reseeding of Multiple-Polynomial Linear Feedback Shift Registers Pattern Generation for a Deterministic BIST Scheme "Compaction of Test Sets Based on Symbolic Fault Simulation" Cellular Automata-Based Pseudorandom Number Generators for Built-In Self-Test "Cost-Effective Generation of Minimal Test Sets for Stuck-at Faults in Combinational Logic Circuits" Accumulator Built-In Self Test for High-Level Synthesis "ROTCO: A Reverse Order Test Compaction Technique" A Multiple Seed Linear Feed-back Shift Register Advanced Automatic Test Generation and Redundancy Identification Techniques Synthesis of Mapping Logic for Generating Transformed Pseudo-Random Patterns for BIST Minimal Test Sets for Combinational Circuits Circuits for Pseudo-Exhaus- tive Test Pattern Generation Test Using Unequiprobable Random Patterns Multiple Distributions for Biased Random Test Patterns Decompression of Test Data Using Variable-Length Seed LFSRs --TR --CTR Sybille Hellebrand , Hua-Guo Liang , Hans-Joachim Wunderlich, A Mixed Mode BIST Scheme Based on Reseeding of Folding Counters, Journal of Electronic Testing: Theory and Applications, v.17 n.3-4, p.341-349, June-August 2001 Hua-Guo Liang , Sybille Hellebrand , Hans-Joachim Wunderlich, Two-Dimensional Test Data Compression for Scan-Based Deterministic BIST, Journal of Electronic Testing: Theory and Applications, v.18 n.2, p.159-170, April 2002 Rainer Dorsch , Hans-Joachim Wunderlich, Reusing Scan Chains for Test Pattern Decompression, Journal of Electronic Testing: Theory and Applications, v.18 n.2, p.231-240, April 2002 Liang Huaguo , Sybille Hellebrand , Hans-Joachim Wunderlich, A mixed-mode BIST scheme based on folding compression, Journal of Computer Science and Technology, v.17 n.2, p.203-212, March 2002

embedded systems;deterministic BIST;BIST;random pattern testing

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Code generation for fixed-point DSPs.

This paper examines the problem of code-generation for Digital Signal Processors (DSPs). We make two major contributions. First, for an important class of DSP architectures, we propose an optimal O(n) algorithm for the tasks of register allocation and instruction scheduling for expression trees. Optimality is guaranteed by sufficient conditions derived from a structural representation of the processor Instruction Set Architecture (ISA). Second, we develop heuristics for the case when basic blocks are Directed Acyclic Graphs (DAGs).

INTRODUCTION Digital Signal Processors (DSPs) are receiving increased attention recently due to their role in the design of modern embedded systems like video cards, cellular telephones and other multimedia and communication devices. DSPs are largely used in systems where general-purpose architectures are not capable of meeting domain specific constraints. In the case of portable devices, for example, the power consumption and cost may make the usage of general-purpose processors prohibitive. The same is true when high-performance arithmetic processing is required to implement dedicated functionality at low cost, as in the case of specific communications Preliminary version of parts of this paper was presented in [Araujo and Malik 1995] at the 1995 ACM/IEEE International Symposiumon System Synthesis, France, September 13-15, 1995, and in [Araujo et al. 1996] at the 1996 ACM/IEEE Design Automation Conference, June 3-7. Author's address: G. Araujo, Institute of Computing, University of Campinas (UNICAMP), Cx.Postal 6176, Campinas, SP, 13081-970, Brazil and S. Malik, Department of Electrical Engineering, Princeton University, Olden St., Princeton, NJ, 08544, USA. Permission to make digital or hard copies of part or all of this work for personal or classroom use is granted without fee provided that copies are not made or distributed for profit or direct commercial advantage and that copies show this notice on the first page or initial screen of a display along with the full citation. Copyrights for components of this work owned by others than ACM must be honored. Abstracting with credit is permitted. To copy otherwise, to republish, to post on servers, to redistribute to lists, or to use any component of this work in other works, requires prior specific permission and/or a fee. Permissions may be requested from Publications Dept, ACM Inc., 1515 Broadway, New York, NY 10036 USA, fax +1 (212) 869-0481, or [email protected]. and computer graphics applications. The increasing usage of these processors has revealed a new set of code generation problems, which are not efficiently handled by traditional compiling techniques. These techniques make implicit assumptions about the regular nature of the target architecture and microarchitecture. This is rarely the case with DSPs, where irregularities in the microarchitecture are the very basis for the efficient computation of specialized functions. Due to hard on-chip memory constraints and hard real-time performance requirements, the code generated for these processors has to meet very high quality standards. Since existing compilation techniques are not up to this task, the vast majority of the code is written directly in assembly language. This research is part of a project directed towards developing compilation techniques that are capable of generating quality code for such processors (http://ee.princeton.edu/spam). The implementation of these techniques forms the compiling infrastructure used in this work, which is called the SPAM compiler. There is a large body of work done in code generation for general purpose pro- cessors. Code generation is, in general, a hard problem. Instruction selection for expressions subsumes Directed Acyclic Graph (DAG) covering, which is an NP-complete problem [Garey and Johnson 1979]. Bruno and Sethi [1976] and Sethi [1975] showed that the problem of optimal code generation for DAGs is NP-complete even for a single register machine. It remains NP-complete for expressions in which no shared term is a subexpression of any other shared term [Aho et al. 1977a]. An efficient solution for a restricted class of DAGs has been proposed in [Prabhala and 1980]. Code generation for expression trees has a number of O(n) solutions, where n is the number of nodes in the tree. These algorithms offer solutions for the case of stack machines [Bruno and Sethi 1975], register machines [Sethi and Ullman 1970] [Aho and Johnson 1976] [Appel and Supowit 1987] and machines with specialized instructions [Aho et al. 1977b]. They form the basis of code generation for single issue, in order execution, general-purpose architectures. The problem of generating code for DSPs and embedded processors has not received much attention though. This was probably due to the small size of the programs running on these architectures, which enabled assembly programming. With the increasing complexity of embedded systems, programming such systems without the support of high-level languages has become impractical. Many of the problems associated to code generation for DSP processors were first brought to light by Lee in [Lee 1988] [Lee 1989], a comprehensive analysis of the architecture features of these processors. Code generation for DSP processors has been studied in the past, but only more recently a number of interesting work has tackled some of its important problems. Marwedel [1993] proposed a tree-based mapping technique for compiling algorithms into microcode architectures. [Liem et al. 1994] uses a tree-based approach for algorithm matching and instruction selection, where registers are organized in classes and register allocation is based on a left-first algo- rithm. Datapath routing techniques have also been proposed [Lanner et al. 1994] to perform efficient register allocation. Wess [1990] proposed the usage of Normal Form Schedule for DSP architectures, and offered a combined approach for register allocation and instruction selection using the concept of trellis diagrams [Wess 1992]. [Kolson et al. 1996] recently proposed an interesting exact approach for register allocation in loops for embedded processors. An overview of the current Code Generation for Fixed-Point DSPs \Delta 3 research work on code generation for DSP processors, and embedded processors in general, can be found in [Marwedel and Goosens 1995]. In this paper, we propose an optimal two phase algorithm which performs instruction selection, register allocation and instruction scheduling for an expression tree in polynomial time, for a class of DSPs. The architecture model here (described in Section 2) is of a programmable highly-encoded Instruction Set Architecture (ISA), fixed-point DSP processor. Formally speaking, this class is an extension of the machine models discussed in Coffman and Sethi [1983]. In the first pass (Section 3), we perform instruction selection and register allocation simultaneously, using the Aho-Johnson algorithm [Aho and Johnson 1976]. The second pass, described in Section 4, is an O(n) algorithm that takes an optimally covered expression tree with n nodes, and schedules instructions such that no memory spills are required. A memory spill is an operation where the contents of a particular register is saved in memory due to a lack of available registers for some operation, and reloaded from memory after that operation is finished. Observe that a memory store operation required by the architecture topology is not considered a memory spill. The proposed algorithm uses the concept of the Register Transfer Graph (RTG) that is a structural representation of the datapath, annotated with ISA information. We show that if the RTG of a machine is acyclic, then optimal code is guaranteed for any program expression tree written for that machine. In this case the DSP is said to have an acyclic datapath. Since DAG code generation is NP-complete, we develop heuristics for the case of acyclic datapaths (Section 5) which again uses the RTG concept. In Section 6 we show the results of applying these ideas to benchmark programs. Section 7 summarizes our major contributions and suggests some open problems. 2. ARCHITECTURAL MODEL DSP processors are irregular architectures, when compared with their general purpose counterparts. This section analyzes the main architecture features which distinguish DSPs from general purpose processors with respect to basic block code generation. It is not the purpose of this section to give a detailed and extensive analysis of these features. A comprehensive analysis of DSP architectures can be found in [Lee 1988][Lee 1989] [Lapsley et al. 1996]. DSPs can be classified according to the type of data they use as fixed-point DSPs and floating-point DSPs. In applications running on a fixed-point DSP, users are responsible for scaling the result of the integer operations. This is automatically done in floating-point DSPs. Floating point units are extremely costly in terms of silicon area and clock cycles. For this reason, a large number of the systems based on DSPs uses fixed-point DSPs. In this case, the acronym DSP will be assumed from now on to mean fixed-point DSP. DSPs have on-chip data memory, based on fast static RAMs and on-chip non-volatile program ROM. Unlike general purpose architectures, DSPs are not designed with cache or virtual memory systems, since data and program streams usually fit into the available on-chip memories. Because on-chip memories are fast and cache misses are not an issue, some DSPs are designed as memory-register architectures [Texas Instruments 1990]. In order to achieve the bandwidth required by its appli- cations, other DSPs architectures provide multiple memory banks [Motorola 1990]. Since performance is an important factor for DSP applications, DSP instructions are usually designed to be fetched in a single machine cycle. In order to achieve that instructions are encoded such as to minimize the number of bits they require. In some architectures [Texas Instruments 1990] this is done by means of data memory pages, where instructions need only to carry the offset of the data within the current page in order to access it. The goal of the design of a DSP datapath is to implement those functional units which can speed up costly operations that frequently occur in the processor application domain. A common example of such units is Multiply and Accumulate (MAC). Due to design requirements, DSP designers frequently constrain the inter-connectivity between registers and functional units. There are two main reasons for this. First, the desired functionality usually requires a particular datapath topol- ogy. Second, broad interconnectivity translates into datapath buses and/or muxes, which results in increased cost and instruction performance degradation. A large number of DSPs are heterogeneous register architectures. These are architectures which contain multiple register files, and instructions that require operands and store the resulting computation in different register files (hence the name het- erogeneous). In general-purpose architectures, instructions usually do not restrict the registers they use, provided they come from the same register file (hence operand registers are homogeneous). This considerably simplifies the code generation prob- lem, since it decouples the tasks of instruction selection from register allocation. Due to this property, we say that general-purpose architectures are homogeneous register architectures. Example 1. An example of a DSP architecture is the TI TMS320C25 Digital Signal Processor (DSP) [Texas Instruments 1990], which will be considered the target architecture for the rest of this paper. This processor is part of the TI TMS320 family of processors, which makes a large number of all commercial DSP processors in use today. The TMS320 family is composed of fixed-point processors (TMS320C1x/C2x/C5x/C54x) which are heterogeneous architectures, and also by a number of floating-point homogeneous architecture DSPs (TMS320C3x/C4x). The TMS320C25 processor contains an ISA with specialized memory-register and register-register instructions. It has three separate register-files (a, p and t) containing a single register each. 3. OPTIMAL INSTRUCTION SELECTION AND REGISTER ALLOCATION In homogeneous register architectures the selection of an instruction has no connection whatsoever with the types of registers that the instruction uses. Selecting instructions for heterogeneous register architectures usually requires allocating register from specific register-files as operands for particular instructions. The strong binding between instruction selection and register allocation indicates that these tasks must be performed together [Araujo and Malik 1995]. Consider, for example, the Intermediate Representation (IR) patterns in Figure 1 corresponding to a subset of the instructions in the TMS320C25 ISA. In Figure 1 each instruction is associated to a tree-pattern whose node is composed of operations (PLUS,MINUS,MUL), registers (a,p,t), constants (CONST) and memory references (m). Code Generation for Fixed-Point DSPs \Delta 5 a m a a a add: apac: mpy: mpyk: p a pac: sacl: a a CONST a a spac: a lac: m lack: CONST Fig. 1. IR patterns for the TMS320C25 processor Instruction Operands Destination Cost Three Address Form add m a,m a 1 a / a +m apac a,p a 1 a / a spac a,p a 1 a / a \Gamma p lack k k a 1 a / k pac p a 1 a / p Table 1. Partial ISA of the TMS320C25 processor These tree-patterns, are represented using three-address form in Table 1. Three-address is a standard compiler representation for instructions, where the destination of the instruction, its two operands (hence the name three address) and the operation it performs are present. Any reference in square brackets is associated to a memory position. Table 1 also lists the cost associated to each instruction. Notice that the instructions implicitly define the registers they use. For example, the instruction apac can only take its operands from registers a and p, and always computes the result back into a. Observe also that operations which transfer data through the datapath like lac m (load register a from memory position m), and pac (move register p into register a) can be represented each as a single node, corresponding to the source register of the transfer operation. The associated cost in this case is only the cost of moving the data from the source register into the destination register. Since registers in DSP architectures are a scarce resource, the final code quality is very sensitive to the cost of routing data through the datapath. 3.1 Problem Definition Optimal instruction selection combined with register allocation is the problem of determining the best cover of an expression tree such that the cost of each pattern match depends not only on the number of cycles of the associated instruction, but also on the number of cycles required to move its operands from the location they 6 \Delta Guido Araujo and Sharad Malik currently are to the location where the instruction requires them to be. 3.2 Problem Solution A solution for this problem is to use a variation of the Aho-Johnson algorithm [Aho and Johnson 1976] such that at each node we keep not only all possible costs for matches at that node, but also all possible costs resulting from matching the node and moving the result from where it is originally computed into any other reachable location in the datapath. Tree-grammar parsers have been used as a way to implement code-generators [Aho et al. 1989] [Fraser et al. 1993] [Tjiang 1993]. They combine dynamic programming and efficient tree-pattern matching algorithms [Hoffman and O'Donnell 1992] for optimal instruction selection. We have implemented combined instruction selection and register allocation using the olive [Tjiang 1993] code-generator generator. olive is based on the techniques proposed in iburg [Fraser et al. 1993]. It takes as input a set of grammar rules where tree-patterns are described in a prefixed linearized format. The IR patterns from Table 1 were converted into the olive description of Figure 2, by rewriting each instruction three-address representation into that format. Notice that the instruction destination registers are now associated to grammar non-terminals, and that these are represented by lower case letters in Figure 2. Fig. 2. Partial olive specification for the TMS320C25 processor (instruction numbers and names on the right are not part of the specification) Rules 1 to 3 and 4 to 5 correspond to instructions that take two operands and store the final result in a particular register (a and p respectively). Rule 6 describes an immediate load into register a. Rules 7 to 10 are associated to data transference instructions and bring the cost of moving data through the datapath into the total cost of a match. We should point out that, for sake of simplicity, we do not represent in Figure 2 all patterns corresponding to commutative operations. For example, instruction add m can be specified in two different ways: PLUS(a,m) and PLUS(m,a). Nevertheless, we will consider for the remainder of this paper that all commutative forms of any operation pattern are available whenever required. If we do not consider instruction scheduling and the associated spills at this point, then the algorithm proposed above is optimal. This follows from the fact that this algorithm is a variation of the provably optimal Aho-Johnson dynamic programming algorithm [Aho and Johnson 1976]. Code Generation for Fixed-Point DSPs \Delta 7 4. SCHEDULING Optimal instruction selection and register allocation for an expression tree is not enough to produce optimal code. For optimal code the instructions must be scheduled in such a way that no memory spills are introduced. Notice that memory positions allocated in the previous phase are not considered spills. They result from the optimal selection of memory-register instructions in the ISA and not from the presence of resource conflicts. Aho and Johnson [1976] showed that, by using dynamic programming, optimal code can be generated in linear time for a wide class of architectures. The schedule they propose is based on their Strong Normal Form Theorem. This theorem guarantees that any optimal code schedule for an expression tree, for a homogeneous register architectural model, can always be transformed into Strong Normal Form (SNF). A code sequence is in SNF if it is formed by a set of code sub-sequences separated by memory storages, where each code sub-sequence is determined by a Strongly Contiguous (SC) schedule. A code sequence is a SC schedule if it is formed as follows: at every selected match m, with child subtrees T 1 and T 2 , continuously schedule the instructions corresponding to subtree T 1 followed by the instructions corresponding to T 2 , and finally the instruction corresponding to pattern m. Wess [1990] used SNF as a heuristic to schedule instructions for the TMS320C25 DSP. 4.1 Problem Definition SC schedules are not an efficient way to schedule instructions for heterogeneous register set architectures. They produce code sequences whose quality is extremely dependent on the order the subtrees are evaluated. Consider for example the IR tree of Figure 3(a). The expression tree was optimally matched using the approach proposed in Section 3 and the target ISA. It takes variables at memory positions m 0 to m 4 and stores the resulting computation into one variable at memory position using m 5 as temporary storage. The code sequences generated for three different schedules and its corresponding three-address representation are showed in Figure 3(b-d). Memory position m 7 was used whenever a spill location was required by the scheduler. For the code of Figure 3(b) the left subtree of each node was scheduled first followed by its right subtree and then the instruction corresponding to the node operation. The opposite approach was used to obtain the code of Figure 3(c). Neither the SC schedules in Figure 3(b) and (c), nor any SC schedule will ever produce optimal code. This is obtained using a non-SC schedule that first schedules the addition then the rest of the tree, as in Figure 3(d). Notice that this schedule is indeed an SNF schedule, since first the subtree corresponding to m 2 is contiguously scheduled followed by a storage operation into memory position m 5 , and by another code sequence resulting from a SC schedule of the rest of the tree. From Figure 3 we can verify how the appropriate SNF schedule minimizes spilling. For example, if the tree of Figure 3(a) is scheduled using left-first, the result of operation \Theta m 1 is first stored in p and then moved into a. Just after that, register a has to be used to route the result of position m 5 . But a still contains a live result (the result of m 0 In this case, the code-generator has to emit code to spill the value of a into memory and recover it (a) a a a a mpy m0 p / t * [m0] lac m3 a / [m3] add m2 a / a pac a / p add m2 a / a sacl m7 [m7] / a sacl m5 [m5] / a lt m1 t / [m1] lac m3 a / [m3] mpy m5 p / t * [m5] mpy m0 p / t * [m0] add m2 a / a sacl m5 [m5] / a pac a / p lt m4 t / [m4] mpy m5 p / t * [m5] mpy m0 p / t * [m0] spac a / a - p lac m7 a / [m7] pac a / p sacl m6 [m6] / a spac a / a - p lt m7 t / [m7] spac a / a - p sacl m6 [m6] / a (b) (c) (d) Fig. 3. (a) Matched IR tree for the TMS320C25; (b) SNF Left-first schedule; (c) SNF Right-first schedule; (d) Optimal schedule later. This would not be required if the scheduler had first stored m 2 into , before loading a with the result of m 0 Problems like the one illustrated above are very common in DSP architectures. The obvious question it raises is - Does there exist a guaranteed SNF schedule such that no spilling is required ? We will prove that this schedule exist, under certain conditions that depend exclusively on the ISA of the target processor. But before doing so, let us define the problem formally: Given an optimally covered expression tree for an heterogeneous register architecture, determine an instruction schedule that does not introduce any spill code. 4.2 Problem Solution This section is divided as follows. In Section 4.2.1 we state and prove a sufficient condition that an heterogeneous register architecture has to satisfy in order to enable spill free schedules. In Section 4.2.2 we introduce the concept of Register Transfer Graph (RTG) and show how it impacts the code generation task. Finally, Code Generation for Fixed-Point DSPs \Delta 9 we prove the existence of an optimal linear time scheduling algorithm for a class DSP architectures which have acyclic RTGs. Let T be an expression tree with unary and binary operations. Let be a function which maps nodes in T to the set R[M , where is a set of N registers, and M the set of memory locations. Let u be the root of an expression tree, with v 1 children of u. Consider that after allocation is performed, registers L(v 1 are assigned to v 1 respectively. and T 2 be the subtrees rooted at v 1 , as in Figure 4. From now on the terms expression tree and allocated expression tree will be used interchangeably, with the context distinguishing if the tree is allocated or not. 4.2.1 Allocation Deadlock Definition 1. An expression tree contains an allocation deadlock iff the following conditions are true: (a) L(v 1 and (c) there exist nodes w 1 and w 2 and w 2 such that L(w 1 The above definition can be visualized in Figure 4. This is the situation when two sibling subtrees T 1 and T 2 contain each at least one node allocated to the same register as the register assigned to the root of the other sibling tree. Using this definition it is possible to propose the following result. Fig. 4. Allocation deadlock in an expression tree Theorem 1. Let T be an expression tree. If T does not have a spill free schedule then it contains at least one subtree which has an allocation deadlock. Proof. Assume that all nodes u in T are such that T u is free of allocation deadlocks and that no valid schedule exist for T . According to Definition 1 T u does not have an allocation deadlock when: (a) In this case, a SNF schedule exists if subtree T 1 is scheduled first followed by subtree This case cannot happen since no non-unary operator of an expression tree takes its two operands simultaneously from the same location. (a) (b) (c) Fig. 5. Trees without allocation deadlock (c) exist for which L(w 2 In this case, it is possible to schedule T 1 first, followed by T 2 and the instruction corresponding to node u. This is a valid schedule because just after the schedule of T 1 is finished only register r 1 is live, and therefore, since no register r 1 exists in , no resource conflict will occur when this subtree is scheduled (Figure 5(a)). exist for which L(w 1 This is symmetric to the previous case. Schedule T 2 first, followed by T 1 and the instruction corresponding to u (Figure 5(b)). exist. This case is trivial, any SC schedule results in a spill free schedule (Figure 5(c)). Since the above conditions can be applied to any node u, T will have a valid schedule that is free of memory spilling code. This contradicts the initial assumption Corollary 1. Let T be an expression tree. If T has no subtree containing an allocation deadlock then it must have a spill free schedule. Moreover this schedule can be computed using the proof of Theorem 1. Proof. Directly from the theorem above. 4.2.2 The RTG Model and Theorem Definition 2. The RTG is a directed labeled graph where each node represents a location in the datapath architecture where data can be stored. Each edge in the RTG from node r i to node r j is labeled after those instructions in the ISA that take operands from location r i and store the result into location r j . The nodes in the RTG represent two types of storage: register files and single registers. Register file nodes represent a set of locations of the same type which can store multiple operands. A datapath single register (or simply single register) is a register file of unitary capacity. Register file nodes are distinguished from single register nodes by means of a double circle. Because of its uniqueness, memory is not described in the RTG. Arrows are used instead to represent memory operations. An incoming (outgoing) arrow pointing to (from) an RTG node r is associated to Code Generation for Fixed-Point DSPs \Delta 11 a load (store) operation from (into) memory. Notice that the RTG is a labeled graph where each edge has labels corresponding to the instructions that require that operation. In other words, if both instructions p and q take one operand in r i and store its result into r j , then the edge from r i to r j will have at least two labels, p and q. We say that an architecture RTG is acyclic if it contains no cycles. As a consequence of that any register transfer cycle in an acyclic RTG has to go through memory 1 . Example 2. Consider, for example, the partial olive description in Figure 2. The RTG of Figure 6 was formed from that description. The numbers in parenthesis on the right side of Figure 2 are used to label each edge of the graph. Not all ISA instructions of the target processor are represented in the description of Figure 2, and therefore not all edges in the RTG of Figure 6 are labeled. Notice that the RTG of the TMS320C25 architecture is acyclic. Other DSP processors also have acyclic RTGs, like the processors TMS320C1X/C2X/C5X and the Fujitsu FDSP-4. This paper proposes a solution for code generation for acyclic RTG architectures. Unfortunately, other known DSPs like the ADSP-2100 and the Motorola 56000 have cyclic RTGs. Nevertheless, as it will be shown later, code generation for these processors can also benefit from the results of this work. a t p1,2,3 Fig. 6. TMS320C25 architecture has an acyclic RTG Theorem 2. If an architecture RTG is acyclic, then for any expression tree there exists a schedule that is free of memory spills. Proof. Let T be an expression tree rooted at u, and v 1 its children, such that and L(v 2 . Let T 1 and T 2 be the subtrees rooted at nodes . Let P k , be subtrees of T with root p k for which the result of operation p k is stored into memory (i.e. L(p k (dark areas in Figure 7) as the subtrees formed in T i after removing all nodes from subtrees P k . We will show that if the RTG is acyclic, an optimal schedule can always be determined by properly ordering the schedules for P k (e.g. P 1 . Here we have to address two cases: (a) Assume that T has no allocation deadlock. Therefore, from Corollary 1 T has an optimal schedule. (b) Now consider that an allocation deadlock is present in T , and that it is caused by registers r 1 and r 2 , as shown in Figure 7. Assume also that there exist paths from r 2 to r 1 in the processor RTG. Observe now that for each node in T 2 (Figure 7) allocated to r 1 , e.g. w 2 , the path that goes from w 2 to its ancestor v 2 (allocated to 1 Observe that a self-loop is not considered an RTG cycle. wp Fig. 7. The RTG Theorem necessarily pass by a node allocated to memory, e.g. p 2 . This comes from the fact that any path from r 1 to r 2 has to traverse memory, given that the RTG is acyclic and that it contains paths from r 2 to r 1 . Notice that one can recursively schedule subtrees P 2 and P 4 in T 2 for which the root was allocated to memory, and that this corresponds to emitting in advance all instructions that store results in . Once this is done, only memory locations are live and the remaining subtree Q 2 contains no instruction that uses r 1 . The nodes that remain to be scheduled are those in subtrees T 1 and . Therefore, the tree T 1 [fug can now be scheduled using Corollary 1 and no spill will be required. Notice that the same result will be obtained if one first recursively schedules all subtrees P 1 (white areas in Figure followed by applying Corollary 1 to schedule subtree Q 1 fug. Based on the proof of Theorem 2 above, an algorithm can be designed which computes the best schedule for an expression tree in any acyclic RTG architecture. We have designed such algorithm and named it OptSchedule. Theorem 3. Algorithm OptSchedule is optimal and has running time O(n), where n is the number of nodes in the subject tree T . Proof. The first part is trivial since OptSchedule implements the proof of Theorem 2. Also from Theorem 2, the algorithm divides T into a set of disjoint subtrees recursively schedules each of them. Therefore, every node in T is visited only once. Hence, the algorithm running time is O(n). Remark 1. If the RTG is acyclic for a particular architecture, then optimal sequential code is guaranteed for any expression tree compiled from programs running on that architecture. Unfortunately, this is not true for those architectures which Code Generation for Fixed-Point DSPs \Delta 13 do not have acyclic RTGs. Nevertheless, expression trees in those architectures can also benefit from this work. Observe from Corollary 1 that if an expression tree is free of allocation deadlocks then it can be optimally scheduled. This is valid for any expression tree generated from any architecture, no matter if this architecture has an acyclic RTG or not. Consider for example that a path is added from p to t in the RTG of Figure 6. This creates a cycle in the architecture RTG, which does not go through memory. On the other hand, any expression tree which do not use this new path is free of allocation deadlocks, and therefore can still be optimally scheduled. Such expression trees could be identified by a simple modification of the instruction selection algorithm. The question of how many of these trees exist in a typical program is still open though. 5. HEURISTIC FOR DAGS Instruction selection for an expression DAG requires DAG covering, which is known to be NP-complete [Garey and Johnson 1979]. In practical solutions to this problem heuristics have been proposed which divide the DAG into its component trees by selecting an appropriate set of trees. However, this dismanteling of the DAG into component trees is not unique and there are several ways in which this can be done. Traditionally, the heuristic employed in the case of homogeneous register architectures is to disconnect multiple fanout nodes of the DAG [Aho et al. 1988]. Dividing a DAG into its component trees requires disconnecting (or breaking) edges in the DAG. For the code generation task, breaking a DAG edge between nodes u and v implies the allocation of temporary storage to save the result of operation u while this is not consumed by operation v. This storage location is traditionally the memory but it can, in general, be any location in the datapath. The key idea proposed here is a heuristic which uses architectural information from the RTG in the selection of component trees of a DAG, such that the resulting code has minimal spills. Consider for example the DAG of Figure 8. Notice that two different approaches can be used to decompose this DAG into its component trees, depending on which edge (e 1 or e 2 is selected to break. From now on, we will represent a broken edge by a line segment traversal to the subject edge. As one can see in Figure 8(b), one extra instruction is generated when the dismanteling heuristic is based on breaking edge e 2 instead of e 1 . Incidentally, the code in Figure 8(a) is also the best sequential code one can generate from the subject DAG. Observe from the architectural description in Table 1, that the multiplication operation requests its operands from memory (m) and t, and that the result of the addition operation always produce its result in the accumulator a. Notice also in Figure 6 that to bring any data from a to register t one has to go through m. From Figure 8 one can see that the result of the addition operation has to be stored into a and must be moved to m or t in order to be used as an operand of the multiplication operation. But to move data from a to t one has to go through memory (m). Suppose the memory position selected to store this temporary result is m 5 . Hence, by breaking DAG edge e 1 one is just assigning in advance a memory operation which will appear on that edge, during the instruction selection phase of the code generation. Notice that the existence of a register-transfer path which always goes through memory whenever data is moved from a to t is a property of the target datapath. Similarly, the register-transfer 14 \Delta Guido Araujo and Sharad Malik lac m2 a / [m2] lac m2 a / [m2] add m3 a / a sacl m5 [m5] / a sacl m5 [m5] / a mpy m5 p / t * [m5] add m4 a / a add m4 a / a mpy m5 p / t * [m5] (a) (b) Fig. 8. (a) Breaking edge e 1 Breaking edge e 2 path from a to p must also pass through memory. Notice also that when edge e 2 is broken, pattern PLUS(a,m) (instruction add m 4 cannot be used to match the addition of m 4 with the result of m 2 in the accumulator a. In this case, instruction lac m 5 in Figure 8(b) has to be issued in order to bring the data from m 5 back to the accumulator adding a new instruction to the final code. a a a a a (1) (1) (1) (2) Fig. 9. Expression DAG after partial register allocation was performed and natural and pseudo-natural edges identified by its corresponding lemma. Code Generation for Fixed-Point DSPs \Delta 15 5.1 Problem Solution The heuristic we propose to address the problem just described is divided into four phases. In the first phase (Section 5.1.1) partial register allocation is done for those datapath operations which can be clearly allocated before any code generation task is performed in the DAG. During the second phase (Section 5.1.2), architectural information is employed to identify special edges in the DAG which can be broken without introducing any loss of optimality for the subsequent tree mapping stages. In the third phase (Section 5.2) edges are marked and disconnected from the DAG. Finally component trees are scheduled and optimal code generated for each component tree (Section 5.2). 5.1.1 Partial Register Allocation. A general property of heterogeneous register architectures is that the results of specific operations are always stored in well defined datapath locations. This does not imply total register allocation because data has to be routed through the datapath to locations required by other instructions. Take for example operations add and mul in the target processor. Notice that they implicitly define the primary storage resources that are used for the operation result. In this case (observe Table 1), no register allocation task is required to determine that registers a and p are respectively used to store the immediate result of operations add and mul. Thus, partial allocation can be performed well in advance, even before the task of breaking the edges of the expression DAG takes place. Again, observe that this is only possible if an operation always uses the same register file to store its immediate result. Consider for example the expression DAG of Figure 9. Notice that partial register allocation can be immediately performed for registers a and p. 5.1.2 Natural Edges. We saw before in Figure 8 that some edges have specific properties originating from the target architecture, which allow us to disconnect them from the DAG without compromising optimality. These edges, termed natural edges, are defined as follows. Definition 3. If the instruction selection matching of edge (u; v) always produces a sequence of data transfer operations in the datapath which pass through memory, edge (u; v) is referred to as a natural edge. (a) (b) r Fig. 10. Natural edges are identified by a single line segment: (a) (u; v) is natural; (b) (u; v) is natural if r i has no self-loop in the RTG Now given an expression DAG D, and a target architecture which has an acyclic RTG. It can be shown that a number of edges in D are natural edges. In order to do that let us state a set of lemmas. Let r 1 and r 2 be a pair of registers in the datapath of an acyclic RTG architecture. Also let , be a function which maps nodes in D into the set of datapath locations R [ M , where R is the set of registers in the datapath and M the set of memory positions. Lemma 1. Let r 1 and r 2 be registers in the architecture RTG, such that there exist no path from r 1 to r 2 . Therefore, any edge (u; v) in D for which and is a natural edge. Proof. Given that a path from registers r 1 to r 2 will be traversed whenever instruction selection is performed on edge (u; v), then a memory operation will always be selected during instruction selection on (u; v), and therefore (u; v) is a natural edge (Figure 10(a)). Lemma 2. Edges (u; v) for which are natural edges only if no self-loop exists on register node r i in the RTG representation of the target architecture (Figure 10(b)). Proof. If an architecture has an acyclic RTG, then any loop in the RTG (which is not a self-loop) will traverse memory. Thus, if register r i has no self-loop in the RTG, then any loop starting at r i will go through memory. Therefore, a memory operation will be selected whenever instruction selection is performed on edge (u; v). Hence (u; v) is a natural edge. Notice that the task of breaking natural edges does not introduce any new operations into the DAG because, as the name implies, during the instruction selection phase a memory operation is naturally selected due to constraints in the architecture datapath topology. As a result, no potential optimality is lost by breaking natural edges. Example 3. Consider each one of the lemmas above and the RTG of Figure 6. Observe the expression DAG of Figure 9 after natural edges have been identified. (1) From Lemma 1 we can see that when r 1 a and r 2 every edge (u; v) such that is a natural edge. (2) Consider now Lemma 2. First take the situation when r From the RTG of Figure 6 observe that register p has no self-loop. Since the RTG is acyclic, then any DAG edge (u; v) such that is a natural edge. Now consider the case when r Register a in Figure 6 contains a self-loop and thus nothing can be said regarding these edges. 5.1.3 Pseudo-Natural Edges. In the following two lemmas we show that DAG edges can sometimes interact such that one edge out of a set of two edges must result in storage in memory. The edges in this set are called pseudo-natural edges. Lemma 3. Consider operation v and its operand nodes u and w in Figure 11(a). If partial register allocation of these operations is such that are (w; v) pseudo-natural edges. Proof. Notice that no binary operation v can take both its operands simultaneously from the same register. We have to consider here two situations: Code Generation for Fixed-Point DSPs \Delta 17 w (a) (b) r r Fig. 11. The selected pseudo-natural edges are identified by a double line segment: (a) One of the edges uses a loop in the RTG; (b) One of the edges goes through memory; (a) If node r i has a self-loop in the architecture RTG, one of the edges, e.g. could be matched by an instruction which takes one operand from r i . On the other hand, when this same instruction matches the other edge, i.e. (w; v), it will make use of a register which is contained in an RTG loop (not a self-loop) that goes from r i back to r i . Similarly as in Lemma 2, matching (w; v) will introduce a sequence of transfer operations which necessarily goes through memory, making (w; v) and (u; v) pseudo-natural edges. (b) If no self-loop node r i exists in the architecture RTG, then both edges are natural edges according to Lemma 2. Lemma 4. Consider operation v and its operand nodes u and w of Figure 11(b). Let the partial register allocation of these nodes be such that . If all RTG paths between each pair of nodes are such that only one path does not go through memory, then (u; v) and (w; v) are pseudo-natural edges. Proof. The proof is trivial and follows from the fact that since operation v cannot take both of its operands from the same register r j at the same time, it has to use two paths in the RTG to bring data from register r j . Since only one path from r j to r i does not go through memory, then the other path has to pass through memory. Based on the lemmas above, we need to decide which edge between (u; v) and (w; v) is to be disconnected from the DAG. Loss of optimality might occur depending on which edge is selected. The selected pseudo-natural edge is identified using a double line segment to distinguish it from natural edges. Unlike natural edges, breaking pseudo-natural edges might result in compromising the optimality of code generation for the component trees. However, there is a good chance that this might not happen in actual practice. Example 4. Consider Lemmas 3 and 4 above and the RTG of Figure 6: Observe the expression DAG of Figure 9 after pseudo-natural edges have been identified. Lemma 3 is satisfied for the case when r (4) In this case, if r only one path exists in the RTG from p to a which does not go through memory. After rules 1-4 of Examples 3 and 4 are applied, the expression DAG of Figure 9 results. Each marked edge in Figure 9 has on its side the number corresponding to a rule used from Examples 3 and 4. 5.2 Dismanteling Algorithm The task of dismanteling an expression DAG may potentially introduce cyclic Read After Write RAW dependencies between the resulting tree components leading to an impossible schedule. A similar problem was also encountered in [Aho et al. 1977a] and [Liao et al. 1995] when the authors studied the problem of scheduling worm- graphs derived from DAGs in single-register architectures. Consider, for example, (b) (a) Fig. 12. (a) Cyclic RAW dependency; (b) Constraining the tree scheduler the reconvergent paths from nodes u to v and the component trees T 1 and T 2 of Figure 12(a). Dismanteling the DAG of Figure 12(a) requires that at least one of the edges of the multiple fanout nodes u and T 2 be disconnected. Assume that edges have been selected as the edges to break. In this case, nodes u, v and tree T 1 can be collapsed into a single component tree T 3 , dismanteling the DAG into trees T 3 and T 4 . When an edge between two nodes is broken, a RAW edge is introduced (dashed lines in Figure 12), in order to guarantee that the original data-dependencies are preserved by the scheduler. In this case, the resulting RAW edges form a cycle between component trees T 3 and T 4 , which results in an infeasible schedule for the component trees. Notice that dismanteling is also possible if edge (T 2 ; w) is broken instead of Figure 12(b)). When this occurs, RAW edge (u; T 2 ) is brought into the resulting component tree (T 3 ). As a consequence, the potential optimality of the tree scheduler algorithm OptSchedule can not be guaranteed anymore, since now it has to satisfy the constraint imposed by the new RAW edge inside T 3 . A possible solution to this problem is to modify the tree scheduler algorithm such that it can satisfy any RAW constraint inserted into the tree. Unfortunately, this is a very difficult task for which an efficient solution seems not to exist. Hence, we have to dismantle the DAG such as to avoid inserting RAW edges into the component trees. From the two situations analyzed above, we can conclude that edges on both reconvergent paths have to be disconnected in order to guarantee proper scheduling Code Generation for Fixed-Point DSPs \Delta 19 of operations inside component trees and between component trees. An algorithm which dismantles the DAG should disconnect edges by using as many natural and pseudo-natural edges as possible. We have designed such an algorithm, which we call Dismantle. The Dismantle algorithm starts by first breaking all natural edges, since breaking these edges adds no cost to the total cost of the final code. After that Dismantle proceeds identifying reconvergent paths. It traverses paths in the DAG looking for edges marked as pseudo-natural edges. If a pseudo-natural edge can be used to break an existing reconvergent path the edge is broken. Otherwise the outgoing edge which starts the reconvergent path at the corresponding multiple fanout node is broken. These edges are marked with a black dot in Figure 13. At this point all reconvergent paths in the expression DAG have been disconnected. Additional edges are then broken such that no node ends up with more than one outgoing edge (these edges are also marked with black dots). The resulting DAG after applying algorithm Dismantle is shown in Figure 13. It decomposes the original DAG into five expression trees Finally, these expression trees are scheduled and code is generated for each expression tree. a a a a a Fig. 13. Resulting component trees after dismanteling 6. EXPERIMENTAL RESULTS DSPstone [Zivojnovic et al. 1994] is a benchmark designed to evaluate the code quality generated by compilers for different DSP processors. DSPstone is divided into three benchmark suites: Application, DSP-kernel and C-kernel. The Application benchmark consists of the program adpcm, a well-known speech encoding algorithm. The DSP-kernel benchmark consists of a number of code fragments, which cover the most often used DSP algorithms. The C-kernel suite aims to test typical C program statements. The DSPstone project was supported by a number of major DSP manufacturers (Analog Devices, AT&T, Motorola, NEC and Texas Instruments). We used this benchmark for experimental evaluations. Scheduling Algorithms Tree Origin Left-first Right-first OptSchedule real update 5 5 5 3 dot product 8 8 8 6 iir one biquad Table 2. Number of cycles to compute expression trees using: Right-Left, Left-Right and OptSchedule 6.1 Expression Trees We have applied algorithm OptSchedule to expression trees extracted from programs in the DSP-kernel benchmark. The metric used to compare the code was the number of cycles that takes to compute the expression tree. Observe from Table 2 that algorithm OptSchedule produces the best code when compared with two SC schedules, what is expected since we have proved its opti- mality. Notice that although SC schedules can sometimes produce optimal code, it can also generate bad quality code, as it is the case for expression tree 6. We can also verify that the same expression tree generates different code quality when different SC schedules are used. The structure of the expression tree dictates the best SC schedule, and this structure is a function of the way the programmer writes the code. 6.2 DAG Types Distribution Expression DAGs were classified in trees, leaf DAGs and full DAGs. Leaf DAGs are DAGs for which only leaf nodes have outdegree greater than one. We classify a DAG as a full DAG if it is neither a tree nor a leaf DAG. As one can see from Table 3, the classification revealed that of all basic blocks analyzed 56% were trees, DSP kernel Basic Blocks Trees Leaf DAGs DAGs real update dot product iir one biquad 1 convolution lms Table 3. Types of DAGs found in typical digital signal processing algorithms Code Generation for Fixed-Point DSPs \Delta 21 DAG DAG Hand-written Standard Dismantle Origin Type Code Heuristic Heuristic complex update F matrix 1x3 L 5 5 0% 5 0% iir one biquad F 15 17 13% convolution lms F 7 9 28% 8 14% Table 4. Experiments with DAGs - Leaf DAG (L); Full DAG 38% leaf DAGs and 6% full DAGs. From the set of benchmarks in Table 3 we have noticed that the majority of the basic blocks found in these programs are trees and leaf DAGs. Another experiment was performed, this time using the DSPstone application benchmark adpcm. As before, basic blocks were analyzed to determine the frequency of trees, leaf DAGs and DAGs. In this case, 94% of the basic blocks in this program were found to be trees, 3% leaf DAGs and 3% full DAGs. Although dynamic counting of basic blocks is required in order to provide information on the impact on execution time, one can reasonably argue that a large portion of this program execution time is spent in processing expression trees. Thus, tree-based code generation is very suitable for this application domain. 6.3 Expression DAGs In Table 4 we list a series of expression DAGs extracted from programs in the DSP-kernel benchmark. We have selected the largest DAG found in each kernel for the purpose of comparison with hand-written code. Hand-written assembly code (or assembly reference code) for each DSP-kernel program is available from the DSPstone benchmark suite [Zivojnovic et al. 1994]. Compiled code was generated for each DAG and the resulting number of cycles for a single loop execution reported in Table 4. Compiled code was also generated using a standard heuristic, which dismantles the DAG by breaking all edges at multiple fanout nodes (column Standard Heuristic). Table 4 shows the number of processor cycles and the overhead with respect to hand-written code. Notice that the overhead is due only to the DAG dismanteling technique. The average overhead when comparing the compiled (Dismantle Heuristic) and the assembly reference code was 7%. Leaf nodes are treated the same way in both heuristics. They are simply duplicated into different nodes - one for each outgoing edge. As a consequence, both heuristics have the same performance for the case of leaf DAGs. The average overhead (Dismantle Heuristic) for the case of full DAGs was higher (11%) than for the case of leaf DAGs (4%). The discrepancy is due to the existence of memory-register and immediate instructions in the processor ISA, which can have zero cost multiple fanout operands when these are memory references or constant values. Although the heuristic gains may seem very small, every byte matters, since DSPs have restricted on-chip memory size, what makes the generation of high 22 \Delta Guido Araujo and Sharad Malik quality code the most important goal for the compiler. 7. CONCLUSION With the increasing demand for wireless and multimedia systems, it is expected that the usage of DSPs will continue to grow. Inspite of this, research on compiling techniques for DSPs has not received the adequate attention. These devices continue to offer new research challenges which originate from the need for high quality code at low cost and power consumption. We have proposed an optimal O(n) instruction selection, register allocation, and instruction scheduling algorithm for expression trees, for a class of heterogeneous register DSP architectures which have acyclic RTGs. We then extend this by proposing heuristics for the case when basic blocks are DAGs. This approach is based on the concept of natural and pseudo-natural edges and seeks to use architectural information to help in the task of dismanteling the expression DAG into a forest of trees. The question on how to generate good code for architectures which have cyclic RTGs remains open though. As it was mentioned before, expression trees generated in these architectures can also benefit from this optimality provided they are free of any allocation deadlock. An interesting question which follows from that is how many of expression trees with this property are generated in programs running on these architectures. More work is under way in order to answer this and other questions. ACKNOWLEDGMENTS This research was supported in part by the Brazilian Council for Research and Development (CNPq) under contract 204033/87-0, and by the Institute of Computing (unicamp), Brazil. --R Code generation using tree matching and dynamic programming. Optimal code generation for expression trees. Code generation for expressions with common subexpressions. Code generation for machineswith multireg- ister operations Generalizations of the Sethi-Ullman algorithm for register allocation Optimal code generation for embedded memory non-homogeneous register architectures Using register-transfer paths in code generation for heterogeneous memory-register architectures Code generation for one-register machine Instructions sets for evaluating arithmetic expressions. Journal of the ACM Engineering a simple Computers and Intractability. Pattern matching in trees. Data routing: a paradigm for efficient data-path synthesis and code generation DSP Processor Fundamentals: Architectures and Features. Programmable DSP architectures: Part I. Programmable DSP architectures: Part II. Instruction selection using binate covering. Marwedel and Goosens Efficient computation of expressions with common subexpressions. Complete register allocation problems. The generation of optimal code for arithmetic expressions. Journal of the ACM 17 Digital Signal Processing Applications with the TMS320 Family. An olive twig. On the optimal code generation for signal flow computation. Automatic instruction code generation based on trellis diagrams. Circuits and Systems --TR Compilers: principles, techniques, and tools Generalization of the Sethi-Ullman algorithm for register allocation Code generation using tree matching and dynamic programming Digital signal processing applications with the TMS320 family (vol. 2) Optimal code generation for embedded memory non-homogeneous register architectures Instruction selection using binate covering for code size optimization Optimal register assignment to loops for embedded code generation Using register-transfer paths in code generation for heterogeneous memory-register architectures Data routing Tree-based mapping of algorithms to predefined structures The Generation of Optimal Code for Arithmetic Expressions The Generation of Optimal Code for Stack Machines Optimal Code Generation for Expression Trees Code Generation for a One-Register Machine Code Generation for Expressions with Common Subexpressions Efficient Computation of Expressions with Common Subexpressions Pattern Matching in Trees Instruction Sets for Evaluating Arithmetic Expressions Code-generation for machines with multiregister operations Code Generation for Embedded Processors Computers and Intractability --CTR Jeonghun Cho , Yunheung Paek , David Whalley, Fast memory bank assignment for fixed-point digital signal processors, ACM Transactions on Design Automation of Electronic Systems (TODAES), v.9 n.1, p.52-74, January 2004 Jeonghun Cho , Yunheung Paek , David Whalley, Efficient register and memory assignment for non-orthogonal architectures via graph coloring and MST algorithms, ACM SIGPLAN Notices, v.37 n.7, July 2002 Alain Pegatoquet , Emmanuel Gresset , Michel Auguin , Luc Bianco, Rapid development of optimized DSP code from a high level description through software estimations, Proceedings of the 36th ACM/IEEE conference on Design automation, p.823-826, June 21-25, 1999, New Orleans, Louisiana, United States Shuvra S. Bhattacharyya , Praveen K. Murthy, The CBP Parameter: A Module Characterization Approach for DSP Software Optimization, Journal of VLSI Signal Processing Systems, v.38 n.2, p.131-146, September 2004 Minwook Ahn , Jooyeon Lee , Yunheung Paek, Optimistic coalescing for heterogeneous register architectures, ACM SIGPLAN Notices, v.42 n.7, July 2007

register allocation;scheduling;code generation

290839

Estimation of lower bounds in scheduling algorithms for high-level synthesis.

To produce efficient design, a high-level synthesis system should be able to analyze a variety of cost-performance tradeoffs. The system can use lower-bound performance estimated methods to identify and puune inferior designs without producint complete designs. We present a lower-bound performance estimate method that is not only faster than existing methods, but also produces better lower bounds. In most cases, the lower bound produced by our algorithm is tight.Scheduling algorithms such as branch-and-bound need fast and effective lower-bound estimate methods, often for a large number of partially scheduled dataflow graphs, to reduce the search space. We extend our method to efficiently estimate completion time of partial schedules. This problem is not addressed by existing methods in the literature. Our lower-bound estimate is shown to by very effective in reducing the size of the search space when used in a branch-and-bound scheduling algorithm.Our methods can handle multicycle operations, pipelined functional units, and chaining of operations. We also present an extension to handle conditional branches. A salient feature of the extended method is its applicability to speculative execution as well as C-select implementation of conditional branches.

Introduction High-level synthesis takes an abstract behavioral specification of a digital system and finds a register-transfer level structure that realizes the given behavior. Usually, there are many different structures that can be used to realize a given behavior. One of the main goals of a synthesis system is to find the structure that best meets the constraints, such as limitations on the number of functional units, registers, power, while minimizing some other parameters like the number of time steps. Operation scheduling and datapath construction are the core of high-level synthesis in obtaining efficient designs in terms of area and speed. Scheduling datapath operations into the best time steps is a task whose importance has been recognized in many systems [12, 14, 15, 16]. Since scheduling is an intractable problem, most high-level synthesis systems use heuristics to find a good schedule. In the absence of good lower-bound estimates, it is difficult to evaluate the performance of heuristics. For a synthesis system to produce efficient designs, it should have the capability to analyze different cost-performance trade-offs. So, a scheduler has to explore the design space with a variety of resource constraints. Instead of producing schedules with each and every resource constraint, a scheduler can use estimation to identify and prune inferior designs. Furthermore, estimation of lower bounds can be used to evaluate a heuristic solution. For an estimation tool to be useful, it has to be much faster than the actual scheduler and the lower bounds it produces should be as tight as possible. We have proposed an efficient estimation technique for the lower-bound performance. We tested our estimation method on a number of benchmarks and compared our results with those of some other known methods in the literature [17, 20]. Our method faster than the methods in [17] and [20]. Our method produces better lower bounds than both of them in many cases. In most cases, the lower bound is tight. Many scheduling algorithms such as the branch and bound methods [3] and multi-schedule methods [2] search through the design space by constructively scheduling operations, one step at a time. During the search process, schedules for a subset of operations in the DFG will be produced and evaluated to check if they can lead to a complete schedule with a target upper-bound performance. Such scheduling methods need a method to estimate lower bounds on the completion time of partial schedules. Since the number of partial schedules is generally very high in a design space search process, this estimation should be faster than the estimation for the entire DFG. In this paper, we have proposed a fast and effective lower-bound estimation method for partially scheduled DFGs. It is an extension of our method for the lower-bound estimation of the entire DFGs. In our approach we defined some very useful data structures that need to be computed only once for a given DFG before the exploration. Using those data structures, our method can compute a lower bound for a partial schedule in O(k) time where k is the number of ready and unfinished operations (defined later in this paper) at the partial schedule. The methods in [5, 17, 20, 21] are originally proposed for estimation for the entire DFG and do not address the estimation from partial schedules. They are too slow to be used for partial schedules. For example, if the methods in [17], [20] are used for partial schedules, they take O(c 2 respectively to compute lower bound for a partial schedule where n is the number of operations to be scheduled and c is the critical path length. We implemented our method and the method in [17] separately into a branch and bound scheduling algorithm and tested on a number of benchmarks. The results show that our method is at least 20 times faster and equally effective in reducing the size of the search space. Our method can be used in any scheduling algorithm that schedules one step at a time. We used this method in an optimal dynamic programming scheduling algorithm [1] we developed, and it drastically reduced the size of the search space. We could obtain optimal schedules in very short computational times. Our methods can handle multi-cycle operations, pipelined functional units and chaining of operations. We extended our methods to handle conditional branches. The extended method is applicable to speculative execution as well as C-select execution of operations in the conditional branches. To our knowledge, no other estimation method in the literature can support speculative execution. In the next section, a brief overview of previous works for lower bound estimation is presented. In section 3, our model and terminology is defined. The method for the estimation of lower bounds for the entire DFG is presented in section 4. The computational complexity of our method as compared to Sharma's method [20] is analyzed in the same section. Our lower-bound estimation method for partial schedules is presented in section 5. Extensions to handle conditional branches and chaining are explained in section 6. Experimental results are presented in section 7 and conclusions are in section 8 . Previous Work There are several methods proposed in the literature for the lower-bound estimation of cost as well as performance. Jain et al [7] proposed a mathematical model for predicting the area-delay curve. Their lower-bound method is very fast, but is too trivial and does not consider precedence constraints at all. The technique proposed by Fernandez and Bussell [4] computes the minimum number of operations which must be scheduled in each sub-interval of time steps. It then derives maximum increase in total execution time over all intervals for not having enough processors to accommodate all the operations in that interval. Their method considers only homogeneous resources and can be applied only to multiprocessor schedules. Their method has been extended to high-level synthesis by Sharma et al [20]. They compute the increase in the length of each interval due to concentration of each type of operations in that interval. They also address lower bounds on area cost including interconnect cost. Their method has a computational complexity of O(nc 2 ) where n is the number of nodes in the DFG and c is the critical path length. The method proposed by Ohm et al [10] estimates lower bounds on functional units as well as registers. Their technique for functional unit estimation is a refinement of the basic technique of [20]. It is not applicable to lower-bound performance estimation. The complexity of their method is O( n(c 2 +n+ e) ) where e is the number of edges in the DFG. The method proposed in [17] uses a relaxation technique of the ILP formulation of the scheduling problem for the lower-bound estimation of performance under resource constraints. It has a computational complexity of O(n is the number of time steps and produced lower bounds as good as [20] on many benchmarks. The method in [21] is similar to [17] in that it relaxes the precedence constraints and solves the relaxed problem using a slack driven list scheduling algorithm. Hu et al [5] proposed a method to estimate lower bounds on iteration time and functional unit cost for functional pipelined DFGs. The complexity of their method is O(nck 2 ) where k is the initiation latency. A recursive technique is proposed in [8] for lower-bound performance estimation and it has a complexity of O(n The complexity of our method for the lower-bound performance estimation of the entire DFGs has a complexity of 3 Model and Definitions A DFG (Data Flow Graph) E) is a (directed acyclic) graph representation of a behavioral description where the set of nodes V represents the set of operations and the set of edges E denotes the set of dependencies (precedence constraints) between the operations. For any two operations should be finished before operation y can start. A node x is called a predecessor of y (and y, a successor of x) if there is a directed path from x to y in G using the arcs in E. An operation without any predecessors is called an input operation and an operation without any successors is called an output operation. Associated with each operation x in V , there is a single type indicating that a functional unit of that type should be used to execute that operation. Resource constraints are given by a set of ordered tuples ! t; D(t) is the delay (number of time steps or clock cycles) an operation of type t takes to complete and N(t) is the number of available resources of type t. If a resource type t is pipelined, one instance of that resource type can be assigned to more than one operation in an overlapping fashion. The latency with which they can overlap is denoted by ffi t . Clearly, Let d x denote the delay of an operation x, which is equal to the delay of the type of resource that executes x. For all (as soon as possible) is the earliest time-step that v can be scheduled to start execution, assuming unlimited resources. For all MSAT (v) (Minimum Steps After This operation) as the minimum number of time steps that any schedule of G is going to take after the completion of operation v, assuming unlimited resources. The critical path length is the minimum number of time-steps that any schedule of G is going to take, assuming unlimited resources. It can be computed as with the maximum taken over all output operations v. 4 Lower-bound estimation for performance The intuitive idea behind our lower-bound estimation is as follows. For each resource type t, we group the operations of that type into three non-overlapping intervals and compute a lower-bound as the sum of the lengths of those intervals. The final lower bound is the maximum among all resource types and all possible groupings of operations of each type. Let P t be the number of operations of type t in the DFG. If there are oe A (i; t) operations of type t with an ASAP value less than or equal to i, then there are at least P operations that cannot be scheduled in the first i time steps. Similarly, if there are oe M (j; t) type t operations with a MSAT value less than j, then there are at least P operations that cannot In this paper, resources refer to single-function functional units only be scheduled to execute in the last j time steps. Thus, there are at least P t operations that cannot either be scheduled in the first i time steps or be scheduled to execute in the last j steps of any schedule. The three intervals considered are : first i steps (interval I 1 ), last steps (interval I 3 ) and an interval between the two (interval I 2 ) that does not overlap with the other two. The lengths of intervals I 1 and I 3 are i and j respectively. The length of I 2 depends on the minimum number of type t operations in that interval as well as the number of available resources of type t. The number of operations in I 2 depends on the ASAP and MSAT values of operations which are determined by data dependencies. Thus, the lower-bound estimation takes into account both precedence constraints and resource constraints. Note that it takes at least i time steps before any of the operations in I 2 can start execution and at least j steps after the last operation has finished execution. The lengths of I 1 and I 3 are independent of the set of operations scheduled in those intervals. Thus, for the purposes of lower-bound estimation, the three intervals are non-overlapping. We denote the minimum number of type t operations in interval I 2 by q(i; j; t) and the minimum length of I 2 due to type t operations by h(i; j; t). As explained above, q(i; j; and the value of h(i; j; t) can be computed from q(i; j; t) as follows. operations can be scheduled into dk=re stages. If type t resources are not pipelined, each stage takes D(t) time steps where D(t) is the delay of a type t resource. If they can be pipelined with a latency ffi t , each stage except the last takes ffi t steps. The last stage in either case takes D(t) steps. Hence, pipelined r e D(t) otherwise. lower bound on the the completion time of any schedule of the given DFG, - is oe A (1; oe A (2; oe A (3; oe A (4; oe A (5; oe A (6; The ordered pair next to each node shows its ASAP and MSAT values Figure 1. An example for the lower-bound estimation for the entire DFG given by max (t;i+j-c) fh(i; c is the critical path length. Proof : The above discussion implies that i is a lower bound for a given i, j and type t. The expression for - is the best lower bound among all i, j and t. The condition makes sure that the intervals are non-overlapping. 2 Figure 1 shows an example for lower-bound computation. The ordered pair next to each node indicates its ASAP and MSAT values respectively. It is assumed that addition takes one time step and multiplication takes two. There are one adder and one non-pipelined multiplier. All values of oe A and oe M for multiplication and addition are shown in the figure. For example, the multiplication operations 1 and 2 each has ASAP value less than or equal to 2. Hence, oe A (2; ) is 2. Similarly, the addition operations 6; 7; 8 and 9 each has MSAT value less than 2. Hence, oe M (2; +) is 4. For the example DFG, the maximum value for - is obtained when i=0 and j=2. Complexity analysis : The ASAP values can be computed in a top-down fashion starting from the input operations as follows. If v is an input operation, ASAP (v) = 1. Otherwise, g. The MSAT values are similarly computed in a bottom-up fashion starting from the output operations as follows. If v is an output operation, g. Let n be the number of operations in the DFG. The number of edges in a DFG will grow linearly in n since the number of inputs of each operation is generally bounded by small number such as two. Hence, the ASAP and MSAT values can be found in O(n) time. The number of resource types is generally bounded by a small number. For any i and t, oe A (i; t) and oe M (i; t) can be found recursively as follows. oe A (i; is the number of type t operations with ASAP value i. Similarly, oe M (i; is the number of type t operations with MSAT value 1. The values for A and M can be found during the computation of ASAP and MSAT values without affecting its complexity. Hence, computing oe A and oe M takes O(c) time where c is the critical path length. Finally, computing lower bound using these values takes O(c 2 ) time because there are O(c 2 ) intervals. Thus, the complexity of our algorithm to estimate lower bound of the entire DFG is O(n The method in [20] is similar to our method in that it estimates the length of each interval of time steps. It computes the required computation cycles of each type as the sum of minimum overlaps of all operations of that type in each interval. Then the difference in the required and available computation cycles of each type is divided by the number of available functional units of that type to get any increase in the length of that interval. In their method, for each interval the minimum overlap of each operation has to be determined. Hence, it has a complexity of O(nc 2 ). Our method computes only the number of operations in each interval in constant time using the precomputed data structures oe A and oe M , thus having a complexity of only O(n 5 Estimating a lower bound for a partially scheduled DFG Scheduling algorithms such as branch and bound methods need to compute lower bounds on the completion times from a large number of partially scheduled DFGs. The methods in [20] or [17] are proposed for the purpose of estimating the lower bound for the entire graph. If those methods are used for estimating lower bounds for partial schedules, the time spent in estimation itself may be so high that the advantage of estimation is nullified. They take O(nc 2 to compute lower bound for each partial schedule. Our method in the previous section also takes time. In this section, we present an extension to that method such that the lower-bound from a partial schedule can be computed more efficiently. Our method takes O(k) time, where k is the number of ready and unfinished operations (defined later in this section) at the partial schedule. In the rest of the paper, we call a partial schedule a configuration. If a configuration R is the result of scheduling r time steps, any unscheduled operation at R can only be scheduled at a time step greater than r. We call r the depth of R denoted by depth(R). Let f(v) denote the time step at which the operation v is scheduled to start execution. An operation x is said to be ready at R if it is not scheduled yet and if all its predecessors are scheduled and finished at R i.e. y such that y is a predecessor of x. Note that d y is the delay of y. The set of ready operations at R is denoted by ready(R). A multi-cycle operation x is said to be unfinished at R if it is scheduled to start execution at a time step less than or equal to depth(R), but not finished at depth(R) i.e. f(x) - depth(R) 1. The set of unfinished operations at R is denoted by unfinished(R). The number of unscheduled operations of type t at configuration R is denoted by unsch(R; t). The basic idea behind the estimation for the partial schedules is as follows. At a partial schedule, a subset of operations is already scheduled so as to satisfy precedence constraints as well as resource constraints. So, instead of considering all possible values for i and j (to divide operations into intervals), we can consider the following special case for each resource type t. For the unscheduled portion of the DFG, we find I = maxfijoe A (i; 0g. Intuitively, I is the number of time steps after the current step before any type t unscheduled operation can start execution. And, J is the number of time steps that any complete schedule from the current configuration takes after the last type t operation has finished executing. The steps in computing a lower bound from a configuration R for a resource type t are as follows. 1. Compute I and J . 2. q(I; J; t) / unsch(R; t). (Since oe A (I; 3. Compute h(I; J; t) from q(I; J; t). (As explained in the previous section) a lower bound on the number of time steps to schedule the remaining operations of type t, the quantity depth(R) J is a lower bound on the completion time of any schedule from R. The maximum of these lower bounds over all resource types t (only if there are any unscheduled operations of type t) gives a lower bound on the completion time of any schedule that configuration R can lead to. The only non-trivial step in computing the lower bound is the computation of I and J (Step 1). The most important merit of our algorithm is that it computes I and J in a very efficient way as described in Figure 2. For each node u and each operation type t, ff(u; t) is defined as the minimum number of time steps that any type t successor of u can start after the starting time of u. The value of ff(u; t) is set to infinity if u has no successors of type t. For each node u and each The value for I is 0 if there is a type t operation in ready(R). Otherwise, it is given by min(a; b) where ff(u; t) and is the number of time steps u finished at R. The value for J is given by min(a; b), where fi(u; t) and is the set of operations in unfinished(R) with a type t successor. Figure 2. Computing the values for I and J operation type t, fi(u; t) is defined as the minimum number of time steps that any schedule of the given DFG is going to take after the completion of all type t successors of u. The value for fi(u; t) can be computed as min v fMSAT (v)g such that v is a type t successor of u. If u has no successors of type t, then : if u is a type t operation, fi(u; otherwise, it is set to infinity. Note that a lower bound for a type t is computed only when there are some unscheduled operations of that type. Therefore, the values of I and J will never be infinity. The formulas for the computation of I and J are based on the following lemma. unscheduled operation x at configuration R, x is either a member of ready(R) or there exists a y in (ready(R) [ unfinished(R)) such that y is a predecessor of x. x be an unscheduled operation. If x is not in ready(R), then there is a predecessor p of x that is not scheduled or scheduled but not finished executing. Among all such p, let q be the farthest from x i.e. the length of the longest path from q to x is maximum among all p. (i) If q is scheduled but not finished executing, q is in unfinished(R). q is not in ready(R), then there is a predecessor q 1 of q that is not scheduled or is in unfinished(R). Note that q 1 is a predecessor of x also. And, q 1 is farther than q from x, which is a contradiction. Hence, q is in ready(R). 2 F or type +;? completion operation in ready(R) completion Figure 3. An example for the estimation of lower-bound completion time of partial schedules Figure 3 shows an example for the estimation of lower-bound completion time of a partially scheduled DFG. It is assumed that the scheduling of the first step is finished. There is one adder and one multiplier both with a delay of one time-step. The lower-bound completion time is 7 time-steps. If the target performance is 6 time-steps, the lower-bound estimation suggests that the selection of operations 2 and 3 in the first time-step is wrong. The method in this section is especially useful for a class of scheduling algorithms that compute the lower bound for a large number of configurations during the design space exploration. The matrices ff and fi for all the operations in the DFG and all resource types are computed only once before the design space exploration. This can be done by computing the transitive closure. For a directed graph, the transitive closure can be computed using depth-first search in O(n(e+n)) where e is the number of edges in the graph [18]. As already explained, e grows linearly in n in a DFG since the number of inputs of each operation is generally bounded by a small number such as two. Thus, ff and fi can be computed in O(n 2 ) time. Note that these values for only the operations in ready(R) and unfinished(R) are used in computing the values of I and J . Hence, a lower-bound for any configuration R can be computed in O(k) time where k is the number of operations in Another major advantage of our method is that it introduces very little memory overhead. The only overhead is to store the matrices ff and fi. When there are a large number of partial schedules the memory requirement is dominated by the amount of information stored at each configuration. Any scheduling algorithm taking full advantage of our lower-bound estimation needs to store very little information at each configuration. 6 Extensions 6.1 Conditional Branches We use the same approach of dividing each type of operations into three non-overlapping intervals. As explained in section 4, the lengths of the first and the third intervals are independent of resource constraints. The length of the second interval is a function of the total resource requirement of the operations that should be scheduled in that interval. If there are no conditional branches, the resource requirement is equal to the number of operations. In the presence of conditional branches, however, more than one operation can share one resource in the same time-step. Effectively, an operation requires only a fraction of the resource. If an operation can share resources with at most /* Computes weights of all operations in the CDFG */ f Partition all the operations into conditional blocks For each operation x in the CDFG f b /\Gamma block of x number of blocks that have a type t operation and mutually exclusive with the block b Figure 4. Outline of the procedure to compute the weights of operations other operations in the same time-step, its minimum resource requirement is 1=(n 1). We refer to this quantity as the weight of that operation. For any given resource type t, the minimum total resource requirement in an interval can be computed as the sum of weights of all type t operations in that interval. Given the weights of individual operations, the computation of the sum of weights of operations in each interval is similar to the computation of the number of operations with no impact on the complexity. For partial schedules, we use the sum of weights of the unscheduled operations of each type t in place of unsch(R; t) at each configuration R. Thus, the only increase in complexity with our extension to conditional branches is due to the computation of the weights of operations. Figure 4 shows an outline of the procedure to compute the weights of individual operations. We partition all the operations into blocks such that all the operations with the same conditional behavior are placed into the same block. Since all the operations in a block have the same control behavior, the concept of mutual exclusiveness between operations can be easily extended to blocks. If x is a type t operation in the block b, then at any given control step x can share a resource with at most one other type t operation from any other block that is mutually exclusive with b. Hence, if there are n blocks that are mutually exclusive with b and that have a type t operation, then the weight of x is 1=(n 1). The method in [19] to handle conditional branches is extension of [20]. In their approach, for each interval the operations from only one conditional path are considered so as to maximize the minimum resource requirement in that interval. Since the conditional path analysis is performed for each interval, their method is very slow. When used in scheduling algorithms for partial schedules, the actual time spent in estimation itself can outweigh the advantage of the resultant pruning. In comparison, our method computes the weights of operations only once and the complexity of the remaining steps remains unchanged. Their method is based on distribute-join representation of CDFGs which is a C-select implementation. In C-select implementation, the operations in conditional branches cannot be executed until the corresponding condition is resolved. Many scheduling algorithms in the recent literature allow execution of branch operations before the corresponding conditional [24, 23, 26, 27]. This is known as speculative execution and is shown to produce faster schedules on many benchmarks [27]. Our estimation method can support both C-select implementation and speculative execution. In C-select implementation, the control precedences are treated the same way as data dependencies are considered in computing ASAP and MSAT values of op- erations. In speculative execution, control dependencies are ignored while computing ASAP and MSAT values. 6.2 Chaining Chaining of operations is handled by dividing time steps into time-units and extending the definitions of MSAT and ASAP values in terms of time-units. Let the length of each time-step be T time-units. Let - v denote the delay of an operation v in terms of time-units. If two operations u and are chained, the functional unit executing u cannot be freed until v is finished [19]. Therefore, if spans across time-steps, this may result in under-utilization of resources. To avoid this, we follow the same assumption as in [19] that an operation v can be chained at the end of operation u only if there is a enough time for v to be finished in the same time step in which u has finished execution. This condition is imposed by the checking that - u mod T 6= 0 and Let A(v) and M(v) be the ASAP and MSAT values of an operation in terms of time-units. The A and M values can be recursively computed similar to the computation of ASAP and MSAT . For any (u; v) 2 E, the earliest time-unit that the execution result of u is available for v is A(u) However, if v cannot be chained to u, v can start execution only at the beginning of the next time step. Hence, A(v) is given by: can be chained to u cannot be chained to u. Similarly, M(v) is given by: ae(v; can be chained to v cannot be chained to v. From the A and M values, the corresponding ASAP and MSAT values are derived. The lower-bound is then computed using the ASAP and MSAT values as explained in section 4. 7 Experimental Results We implemented our methods in C language on a SUN Sparc-2 workstation. We tested them using a number of benchmarks in the literature. The benchmarks we used are the AR Filter [6], the fifth-order elliptic Wave Filter [13], twice unfolded Wave Filter, the complex Biquad recursive digital Filter [13], the sixth-order elliptic Bandpass Filter [13], Discrete Cosine Transform [11], and Fast Discrete Cosine Transform [9]. For the Biquad Filter example, we used three time steps for the multiplier and one for adder (we used the same resource type adder to do addition, subtraction and comparison). For all the other examples, we used two time steps for multiplication and one for adder. 7.1 Lower bound estimation of partial schedules in branch and bound methods As mentioned in section 1 , branch and bound scheduling methods rely on estimating lower bounds for partial schedules to keep the design space from exploding. Generally, lower bounds need to be estimated for a large number of partial schedules (configurations). If the time spent in lower-bound estimation itself is too high, it will have a big negative impact on the over-all time taken by the scheduling algorithm. We implemented a branch and bound scheduling algorithm [3] and tested on the benchmarks. We first find a schedule using a list scheduling algorithm. We use that performance as an upper bound in the branch and bound algorithm and search the design space exhaustively for an optimum schedule. From each partial schedule, we estimate a lower bound for the schedule completion time. If it exceeds the upper-bound, the partial schedule cannot lead to a complete schedule with the target performance and it is not explored further. We separately measured the time spent in lower-bound estimation using our method of section 5 and Rim's method [17]. The results are reported in table 1. Our method is at least 20 times faster in all cases. As a measure of the effectiveness of the lower-bound estimation in reducing the size of the search space, we also measured the number of configurations visited using each method separately. These results are also reported in table 1. Both the methods are equally effective. Since their estimation is very slow, the CPU time taken using their method is more than the time taken without using any lower-bound estimation in a few cases. However, in a majority of the cases, the search space exploded without lower-bound estimation, thus showing the necessity of estimation in Resources Configurations Our method Rim [17] Our method Rim AR Filter 2 3 0.35 8.2 1769 3497 Once unrolled 2 3 0.13 2.9 770 723 Twice unrolled Filter 2 3 4.0 81.2 33182 27767 Filter Fast Discrete 1 1 0.14 3.1 688 892 Table 1. CPU time and number of configurations in branch and bound algorithm branch and bound scheduling algorithms. Our method is more suitable than the existing methods to be used in such scheduling algorithms. We incorporated this method in a dynamic programming scheduling (DPS) algorithm [1] that we developed and obtained excellent results. 7.2 Lower-bound estimation for the entire DFGs We tested all the benchmarks with different resource constraints for pipelined multiplier (latency is 1) and non-pipelined multiplier. For all the cases, our lower bound is compared with an optimal solution we obtained using our DPS algorithm [1]. In tables 2 and 3, we present the lower bounds for some of the cases as obtained by our method of section 4 . The column DPS in the tables shows the number of steps in an optimal solution. The lower bound is tight in 156 out of 198 cases. In 22 more cases, the difference is only one step. We also implemented the algorithms by Rim [17] and Sharma [20], and compared the results with ours. Our method gives better lower bound than [17] in nine cases. In five cases, our lower bounds are better than [20]. In [17], the lower bounds by one more method, Jain [7] are also reported. Those are copied into the second last column (Jain) of our our lower bound and an Resources optimum solution other lower bounds lower bound difference Rim [17] Jain [7] Sharma [20] AR Filter 1 3 Twice unrolled Wave Filter 3 2 50 50 0 50 Fast (*) Complex multiplication takes 3 time steps (y) Our lower bound for this case is better than Rim's (z) Our lower bound for this case is better than Sharma's Table 2. Lower bounds with non-pipelined multiplier tables. For the benchmarks and cases not reported in our tables, our lower bounds are identical to Rims' [17]. The average CPU times are 21 ms, 25 ms and 270 ms for our method, Rim's method and Sharma's method respectively. Thus, our method is faster than the fastest non-trivial method in the literature [17] and produces better lower bounds in more cases. Our method is one order faster than the method in [20] and still produces better results in some cases. The lower bounds of [7] are far inferior to ours. our lower bound and an Resources optimum solution other lower bounds lower bound difference Rim [17] Jain [7] Sharma [20] Fast 1 1 26 26 0 26 - 26 Transform (*) Complex multiplication takes 3 time steps Table 3. Lower bounds with pipelined multiplier 7.3 Results for CDFGs Table 4 shows the results for examples with conditional behavior - Maha from [14], Parker from High Level Synthesis Benchmark Suite, Kim from [25], Waka from [22] and MulT from [23]. The resources column lists the number of adders, subtracters and comparators used in each case. All additions, subtractions and comparisons are single-cycle. We presented the number of time steps in the schedules obtained by our DPS algorithm. The lower bounds in all but a few cases are tight. We obtained schedules both with C-select implementation and by allowing speculative execution. In C-select implementation, operations from mutually exclusive branches can always share resources. However, since control precedences are strict precedences, critical path length may increase. In table 4, Maha and Parker are two examples with a high degree of branching. In the C-select imple- mentation, the advantage of conditional resource sharing is nullified by the increase in the critical path length. The length of the schedules could not be reduced even by adding more resources. In comparison, speculative execution gives much superior results and adding more resources reduces schedules lengths. Benchmark Resources # Time Steps C-select Spec. Exec. Maha 1,1,1 11 y7 Maha 2,1,1 11 y6 Maha 2,2,2 11 5 Parker 1,1,1 11 y6 Parker 2,2,1 11 5 Parker 2,2,2 11 5 y For these cases, lower bound is one step less. All other lower bounds are tight. Table 4. C-select and Speculative execution in Conditional Branch benchmarks 8 Conclusions and future research We have presented simple and efficient techniques for estimating lower-bound completion time for the scheduling problem. The proposed techniques can handle multi-cycle operations, pipelined functional units, conditional branches and chaining of operations. Our method for the entire DFGs is faster and produces better lower bounds than [17] and [20]. We have also presented an extension to our technique that is especially suitable for finding lower-bound for partially scheduled DFGs. The extended method is very useful to keep the search space from exploding in scheduling algorithms such as branch and bound method. Exising methods in the literature do not give any special consideration for computing the lower bounds for partial schedules. We conducted extensive experiment using our method and the fastest non-trivial method known in the literature [17] for the estimation of partial schedules in a branch and bound algorithm. Our method is found to be at least 20 times fatser than theirs while being equally effective in reducing the size of the search space. We are currently investigating estimation of lower bounds in the presence of loops and when multi-function functional units are used. We are also investigating estimation of lower bounds with additional constraints such as interconnect and storage. --R "Optimum Dynamic Programming Scheduling under Resource Constraints" "A Multi-Schedule Approach to High-Level Synthesis" "Some Experiments in Local Microcode Compaction for Horizontal Machines" "Bounds on the Number of processors and Time for Multi-processor Optimal Schedules" "Lower Bounds on the Iteration Time and the Number of Resources for Functional Pipelined Data Flow Graphs" "Experience with the ADAM Synthesis System" "Predicting system-level area and delay for pipelined and non-pipelined designs" "A Recursive Technique for Computing Lower-Bound Performance of Schedules" "A new approach to pipeline optimization" "Comprehensive Lower Bound Estimation from Behavioral Descriptions" "Personal Communication" "Slicer: A State Synthesizer for Intelligent Silicon Compiler" "A High Level Synthesis Technique Based on Linear Pro- gramming" "MAHA: A Program for Datapath Synthesis" "SEHWA: A Software Package for Synthesis of Pipelines for Synthesis of Pipelines from Behavioral Specifications" "Lower-Bound Performance Estimation for the High-Level Synthesis Scheduling Problem" "Algorithms in C" "Estimation and Design Algorithms for the Behavioral Synthesis of ASICS" "Estimating Architectural Resources and Performance for High-Level Synthesis Applications" "Estimating Implementation Bounds for Real Time DSP Application Specific Circuits" "A resource sharing and control synthesis method for conditional branches" "Global Scheduling Independent of Control Dependencies Based on Condition Vectors" "A Tree-Based Scheduling Algorithm For Control-Dominated Circuits" "A Scheduling Algorithm For Conditional Resource Sharing" "Global Scheduling For High-Level Synthesis Applications" "A New Symbolic Technique for Control-Dependent Scheduling" --TR Experience with ADAM synthesis system Algorithms in C Global scheduling independent of control dependencies based on condition vectors A tree-based scheduling algorithm for control-dominated circuits Comprehensive lower bound estimation from behavioral descriptions Global scheduling for high-level synthesis applications MAHA A Multi-Schedule Approach to High-Level Synthesis Estimation and design algorithms for the behavioral synthesis of asics A new approach to pipeline optimisation --CTR Shen Zhaoxuan , Jong Ching Chuen, Lower bound estimation of hardware resources for scheduling in high-level synthesis, Journal of Computer Science and Technology, v.17 n.6, p.718-730, November 2002 Helvio P. Peixoto , Margarida F. Jacome, A new technique for estimating lower bounds on latency for high level synthesis, Proceedings of the 10th Great Lakes symposium on VLSI, p.129-132, March 02-04, 2000, Chicago, Illinois, United States Margarida F. Jacome , Gustavo de Veciana, Lower bound on latency for VLIW ASIP datapaths, Proceedings of the 1999 IEEE/ACM international conference on Computer-aided design, p.261-269, November 07-11, 1999, San Jose, California, United States Margarida F. Jacome , Gustavo de Veciana, Lower bound on latency for VLIW ASIP datapaths, Readings in hardware/software co-design, Kluwer Academic Publishers, Norwell, MA, 2001 Margarida F. Jacome , Gustavo de Veciana , Viktor Lapinskii, Exploring performance tradeoffs for clustered VLIW ASIPs, Proceedings of the 2000 IEEE/ACM international conference on Computer-aided design, November 05-09, 2000, San Jose, California

high-level synthesis;lower-bound estimated;dynamic programming;scheduling

291058

Value speculation scheduling for high performance processors.

Recent research in value prediction shows a surprising amount of predictability for the values produced by register-writing instructions. Several hardware based value predictor designs have been proposed to exploit this predictability by eliminating flow dependencies for highly predictable values. This paper proposed a hardware and software based scheme for value speculation scheduling (VSS). Static VLIW scheduling techniques are used to speculate value dependent instructions by scheduling them above the instructions whose results they are dependent on. Prediction hardware is used to provide value predictions for allowing the execution of speculated instructions to continue. In the case of miss-predicted values, control flow is redirected to patch-up code so that execution can proceed with the correct results. In this paper, experiments in VSS for load operations in the SPECint95 benchmarks are performed. Speedup of up to 17% has been shown for using VSS. Empirical results on the value predictability of loads, based on value profiling data, are also provided.

INTRODUCTION Modern microprocessors extract instruction level parallelism (ILP) by using branch prediction to break control dependencies and by using dynamic memory disambiguation to resolve memory dependencies [1]. However, current techniques for extracting ILP are still insufficient. Recent research has focused on value prediction hardware for dynamically eliminating flow dependencies (also called true dependencies) [2], [3], [4], [6], [7], [8], [9]. Results have shown that values produced by register-writing instructions are potentially highly predictable using various value predictors: last-value, stride, context-based, two-level, or hybrid predictors. This work illustrates that value speculation in future high performance processors will be useful for breaking flow dependencies, thereby exposing more ILP. This paper examines ISA, hardware and compiler synergies for exploiting value speculation. Results indicate that this synergy enhances performance on difficult, integer benchmarks. Prior work in value speculation utilizes hardware-only schemes (e.g. [2], [3]). In these schemes, the instruction address (PC) of a register-writing instruction is sent to a value predictor to index a prediction table at the beginning of the fetch stage. The prediction is generated during the fetch and dispatch stages, then forwarded to dependent instructions prior to their execution stages. A value- speculative dependent instruction must remain in a reservation station (even while its own execution continues), and be prevented from retiring, until verification of its predicted value. The predicted value is compared with the actual result at the state-update stage. If the prediction is correct, dependent instructions can then release reservation stations, update system states, and retire. If the predicted value is incorrect, dependent instructions need to re-execute with the correct value. Figure 1 Figure 1. Pipeline Stages of Hardware Value Speculation Mechanism for Flow Dependent Instructions. The dependent instruction executes with the predicted value in the same cycle as the predicted instruction. illustrates the pipeline stages for value speculation utilizing a hardware scheme. Little work has been done on software-based schemes to perform value prediction and value speculation of dependent instructions. In a related approach to a different problem, the memory conflict buffer [1] was presented to dynamically disambiguate memory dependencies. This allows the compiler to speculatively schedule memory references above other, possibly dependent, memory instructions. Patch-up code, generated by the compiler, ensures correct program execution even when the memory dependencies actually occur. Speculatively scheduled memory references improves performance by aggressively scheduling references that are highly likely to be independent of each other. Likewise, value-speculative scheduling attempts to improve performance by aggressively scheduling flow dependencies that are highly likely to be eliminated through value prediction. Patch-up code is used when values are miss-predicted. We apply this scheme to value speculation and propose a combined hardware and software solution, which we call value speculation scheduling (VSS). Hardware pipeline stages for the VSS scheme are shown in Figure 2. Two new instructions, LDPRED and UDPRED, are introduced to interface with the value predictor during the execution stage. LDPRED loads the predicted value generated by the predictor into a specified general-purpose register. UDPRED updates the value predictor with the actual result, resetting the device for future predictions after a miss-prediction. Figure 3 shows an example of using LDPRED and UDPRED to perform VSS. In the original code sequence of Figure 3(a), instructions I1 to I6 form a long flow dependence chain, which must execute sequentially. If the flow dependence from Figure 2. Pipeline Stages of Value Speculation Scheduling Scheme. Two new instructions, LDPRED and UDPRED, interface with the value predictor during the execution stage. (a) Original code I3: LW R4 # 0(R3) I4: ADD R5 # R4, 1 I5: Next: . (b) New code after value speculation of R4 (predicted instruction I3) I3: LW R4 # 0(R3) I7: LDPRED R8 # index // load prediction into R8 I5': I8: BNE Patchup R8, R4 // verify prediction Next: . Patchup: I9: UDPRED R4, index // update predictor with R4 I4: ADD R5 # R4, 1 I5: I10: JMP Next Figure 3: Example of Value Speculation Scheduling. instruction I3 to I4 is broken, via VSS, the dependence height of the resulting dependence chain is shortened. Furthermore, ILP is exposed by the resulting data dependence graph. Figure 4 shows the data dependence graphs for the code sequence of Figure 3 before and after breaking the flow dependence from instruction I3 to I4. Assume that the latencies of arithmetic, logical, branch, store, LDPRED and UDPRED instructions are 1 cycle, and that the latency of load instructions is 2 cycles. Then, the schedule length of the original code sequence of Figure 4(a), instructions I1 to I6, is seven cycles. By breaking the flow dependence from instruction I3 to I4, VSS results in a schedule length of five cycles. Figure 4(b) illustrates the schedule now possible due to reduced overall dependence height and ILP exposed in the new data dependence graph. Fetch Dispatch Execute State- Update Value Predictor Prediction Verification Fetch Dispatch Execute State- Update Predicted Value Actual Value (Predicted (Dependent Fetch Dispatch Execute State- Update Value Predicted Value LDPRED UDPRED This improved schedule length, from seven cycles to five cycles, does not consider the penalty associated with miss- prediction due to the required execution of patch-up code. The impact of patch-up code on performance will be discussed in section 3. Figure 4. Data Dependence Graphs for Codes of Figure 3. The numbers along each edge represent the latency of each instruction. In 4(a), the schedule length is seven cycles. In 4(b), because of exposed ILP and dependence height reduction, the schedule length is reduced to five cycles. In Figure 3(b), the value speculation scheduler breaks the flow dependence from instruction I3 to I4. Instructions I4, I5 and I6 now form a separate dependence chain, allowing their execution to be speculated during scheduling. They become instructions I4' I5' and I6', respectively. An operand of instruction I4' is modified from R4 to R8. Register R8 contains the value prediction for destination register R4 of the predicted instruction I3. Instruction I7, LDPRED, loads the value prediction for instruction I3 into register R8. When the prediction is incorrect (R8-R4), instruction I9, UDPRED, updates the value predictor with the actual result of the predicted instruction, from register R4. Note that the resulting UDPRED instruction is part of patch-up code and its execution is only required when a value is miss-predicted. To ensure correct program execution, the compiler inserts the branch instruction, I8, after the store instruction, I6', to branch to the patch-up code when the predicted value does not equal the actual value. The patch-up code contains UDPRED and the original dependent instructions, I4, I5 and I6. After executing patch-up code, the program jumps to the next instruction after I8 and execution proceeds as normal. Each LDPRED and UDPRED instruction pair that corresponds to the same value prediction uses the same table entry index into the value predictor. Each index is assigned by the compiler to avoid unnecessary conflicts inside the value predictor. While the number of table entries is limited, possible conflicts are deterministic and can be factored into choosing which values to predict in a compiler approach. A value predictor design, featuring the new LDPRED and UDPRED instructions, will be described in section 2. By combining hardware and compiler techniques, the strengths of both dynamic and static techniques for exploiting ILP can be leveraged. We see several possible advantages to VSS: . Static scheduling provides a larger scheduling scope for exploiting ILP transformations, identifying long dependence chains suitable for value prediction and then re-ordering code aggressively. . Value-speculative dependent instructions can execute as early as possible before the predicted instruction that they depend. . The compiler controls the number of predicted values and assigns different indices to them for accessing the prediction table. Only instructions that the compiler deems are good candidates for predictions are then predicted, reducing conflicts for the hardware. . Patch-up code is automatically generated, reducing the need for elaborate hardware recovery techniques. . Instead of relying on statically predicted values (e.g., from profile data), LDPRED and UDPRED access dynamic prediction hardware for enhanced prediction accuracy. . VSS can be applied to dynamically-scheduled processors, statically-scheduled (VLIW) processors, or EPIC (explicitly parallel instruction processors [14]. There is a drawback to VSS. Because static scheduling techniques are employed, value-speculative instructions are committed to be speculative and therefore always require predicted values. Hardware only schemes can dynamically decide when it is appropriate to speculatively execute instructions. The dynamic decision is based on the value predictor's confidence in the predicted value, avoiding miss-prediction penalty for low confidence predictions. The remainder of this paper is organized as follows: Section 2 examines the value predictor design for value speculation scheduling. Section 3 introduces the VSS algorithm. Section 4 presents experimental results of VSS. Section 5 concludes the paper and mentions future work. 2. VALUE PREDICTOR DESIGN Microarchitectural support for value speculation scheduling (VSS) is in the form of special-purpose value predictor hardware. Value prediction accuracy directly relates to performance improvements for VSS. Various value predictors, such as last-value, stride, context-based, two- level, and hybrid predictors [2], [3], [4], [6], [7], [9], (a) Before breaking dependence (b) After breaking dependence from I3 to I4 provide different prediction accuracy. Value predictors with the most design complexity, in general, provide for the highest prediction accuracy. In order to feature LDPRED and UDPRED instructions for VSS, previously proposed value predictors must be re-designed slightly. Figure 5 shows the block diagram of a value predictor that includes LDPRED and UDPRED instructions. In this value predictor, there are three fundamental units, the current state block, the old state block and the prediction hardware block. The current state block may contain register values, finite state machines, history information, or machine flags, depending on the prediction method employed. The old state block hardware is a duplicate of the current state block hardware. Predictions are generated by the prediction hardware with input from the current state block. Various prediction mechanisms can be used. For example, generating the prediction as the last value (last value predictors [2], [3]). Or, generating the prediction as the sum of the last value and the stride, which is the difference between the most recent last values (stride predictors [4], [6], [7], [9]). Also, two-level predictors [7] allow for the prediction of recently computed values. For two-level predictors, a value history pattern indexes a pattern history table, which in turn is used to index a value prediction from recently computed values. Two-level value prediction hardware is based on two-level branch prediction hardware. Figure 5. Block Diagram of Value Predictor featuring LDPRED and UDPRED. Both the LDPRED and UDPRED instructions contain an immediate operand that specifies the value predictor table index. In general (independent of the prediction hardware chosen) the LDPRED instruction performs three actions. The compiler assigned number indexes each action. First, the prediction hardware generates the predicted value by using input from the current state block. Second, current state information is shifted to the old state block. Last, the current state block is updated based on the predicted value from the prediction hardware. Information used by the prediction hardware is updated simultaneously with the current state block update. Note that for the LDPRED instruction, the predicted value is used to update the current state block speculatively. The compiler assigned number also indexes the operation of the UDPRED instruction. When the value prediction is incorrect, the patch-up basic block of Figure 3(b) must be executed. The execution of UDPRED instructions only occurs in patch-up code, or only when values are miss- predicted. The UDPRED instruction causes the update of both the current state block and the prediction hardware with the actual computed value and the old state block. If the compiler can ensure that each LDPRED/UDPRED instruction pair is executed in turn (each prediction is verified and value predictions are not nested) the old state block requires only one table entry. The same table entry in the old state block is updated by every LDPRED instruction, and used by every UDPRED instruction, in the case of miss-prediction. Figure 6. Hybrid Predictor (Stride and Two-Level). Saturating counters are compared to select between the prediction techniques. In the VSS scheme, a prediction needs to be generated for each LDPRED instruction. There is no flag in the value predictor to indicate if a value prediction is valid or not. The goal of the value predictor is to generate as many correct predictions as possible. In this paper, stride, two-level and hybrid value predictors [7] are implemented to find the design which provides the highest prediction accuracy for use in the VSS scheme. Stride predictors predict arrays and loop induction variables well. Two-level predictors capture the recurrence of recently used values and generate predictions based on previous patterns of values. However, neither of them alone can obtain high prediction accuracy for all programs, which exhibit different characteristics. Therefore, hybrid value predictors, consisting of both stride and two-level prediction are designed to cover both of these situations. Current State Old State Prediction Hardware Actual Value (LDPRED, Predicted Value LDPRED UDPRED Prediction Index Stride Two-Level Counters for Counters for Two-Level Figure 6 shows such a hybrid predictor that obtains high prediction accuracy. The selection between the stride predictor and the two-level predictor is different from that in [7]. Every table entry has a saturating counter in the stride predictor and in the two-level predictor. The saturating counter increments when its corresponding prediction is correct, and decrements when its prediction is incorrect. Both saturating counters and predictors are updated for each prediction, regardless of which prediction is actually selected. The hybrid predictor selects the predictor with the maximum saturating counter value. In the event of a tie, the hybrid predictor favors the prediction from the two-level predictor. Prediction accuracy results for the three value predictors will be presented in section 4. 3. VALUE SPECULATION SCHEDULING Performance improvement for value speculation scheduling (VSS) is affected by prediction accuracy, the number of saved cycles (from schedule length reduction) and the number of penalty cycles (from execution of patch-up code). Suppose that after breaking a flow dependence, value-speculative dependent instructions are speculated, saving S cycles in overall schedule length when the prediction is correct. Patch-up code is also generated and requires P cycles. Prediction accuracy for the speculated value is X. In this case, speedup will be positive if S > (1- holds. For the example of Figure 3(b) VSS saves 2 cycles (from 7 cycles to 5 cycles) and the resulting patch-up code contains 5 instructions, requiring 3 cycles in an ILP processor. Therefore, for positive speedup, the prediction accuracy must be greater than 33%. If the actual prediction accuracy is less, performance will be degraded by VSS. With these performance considerations in mind, an algorithm for VSS is proposed in Figure 7. The first step is to perform value profiling. The scheduler must select highly predictable instructions to improve performance through VSS. Results from value profiling under different inputs and parameters have been shown to be strongly correlated [5], [6]. Therefore, value profiling can be used to select highly predictable instructions on which to perform value speculation. Value profiling can be performed for all register-writing instructions. If profiling overhead is a concern, a filter may be used to perform value profiling only on select instructions. Select instructions may be those that reside in critical paths (long dependence height) or those that have long latency (e.g., load instructions). In [5], estimating and convergent profiling are proposed to reduce profiling overhead for determining the invariance of instructions. Similar techniques could be applied for determining the value predictability of instructions. Next, the value speculation scheduler performs region formation. Treegion formation [10] is the region type chosen for our experiments. A treegion is a non-linear region that includes multiple execution paths in the form of a tree of basic blocks. The larger scheduling scope of treegions allows the scheduler to perform aggressive control and value speculation. A data dependence graph is then constructed for each treegion. In step four, a threshold of prediction accuracy is used to determine whether or not to perform value speculation on each instruction. For each instruction, the scheduler queries the value profiling information to get the estimate of its predictability. If the predictability estimate is greater than the threshold, value prediction is performed. For aggressive scheduling, more instructions can be speculated by choosing a low threshold. Suggested values for the threshold are derived from experimental results in section 4. When an instruction is selected for value prediction, a LDPRED instruction is inserted directly after it. The LDPRED instruction has an immediate value that is assigned by the scheduler to be its chosen index into the value predictor. A new register is also assigned as the destination of the LDPRED instruction. Once the new destination register has been chosen for the LDPRED instruction, any dependent instruction(s) need to update their source register(s) to reflect the new dependence on the LDPRED instruction. Only the first dependent instruction in a chain of dependent instructions needs to update its register source, the remaining dependencies in the chain are 1. Perform Value Profiling 2. Perform Region Formation 3. Build Data Dependence Graph for Region 4. Select Instruction with Prediction Accuracy (based on Value Profiling) greater than a Threshold 5. Insert LDPRED after Predicted Instruction (selected instruction of step 6. Change Source Operand of Dependent Instruction(s) to Destination Register of LDPRED 7. Insert Branch to Patch-up Code 8. Generate Patch-up Code (which contains UDPRED) 9. Repeat Steps 4 - 8 until no more Candidates Found 10. Update Data Dependence Graph for Region 11. Perform Region Scheduling 12. Repeat Steps 2 - 11 for each Region Figure 7. Algorithm of Value Speculation Scheduling. unaffected. Even though more than one chain of dependent instructions may result from just one value prediction, only one LDPRED instruction is needed for each value prediction. In step seven, a branch to patch-up code is inserted for repairing miss-predictions. Only one branch per data value prediction is required and the scheduler determines where this branch is inserted. Once the location of the branch is set, all instructions in all dependence chains between the predicted instruction and the branch to patch-up code are candidates for value-speculative execution. It is therefore desirable to schedule any of these instructions above the predicted instruction. Actual hardware resources will restrict the ability to speculatively execute these candidates for value speculation. Also, as all candidates for value speculation are duplicated in patch-up code, their number directly affects the penalty for miss-prediction. These factors affect the scheduler's decision on where to place the branch to patch-up code. In step eight, patch-up code is created for repairing miss- predictions. The patch-up code contains the UDPRED instruction, a copy of each candidate for value-speculative execution, and an unconditional jump back to the instruction following the branch to patch-up code. The UDPRED instruction uses the same immediate value, assigned by the scheduler, as its corresponding LDPRED instruction for indexing the value predictor. The other source operand for the UDPRED instruction is the destination register of the predicted instruction (the actual result of the predicted instruction). The UDPRED instruction index and the actual result are used to update the value predictor. Finally, in steps ten and eleven, the data dependence graph is updated to reflect the changes and treegion scheduling is performed. Because of the machine resource restrictions and dependencies, not all candidates for value speculation are speculated above the predicted instruction. Section 4 shows the results of using different threshold values for determining when to do value speculation. 4. EXPERIMENTAL RESULTS The SPECint95 benchmark suite is used in the experiments. All programs are compiled with classic optimizations by the IMPACT compiler from the University of Illinois [11] and converted to the Rebel textual intermediate representation by the Elcor compiler from Hewlett-Packard Laboratories [12]. Then, the LEGO compiler, a research compiler developed at North Carolina State University, is used to insert profiling code, form treegions, and schedule instructions [10]. After instrumentation for value profiling, intermediate code from the LEGO compiler is converted to C code. Executing the resultant C code generates value profiling data. For the experiments in value speculation scheduling (VSS), load instructions are filtered as targets for value speculation. Load instructions are selected because they are usually in critical paths and have long latencies. Value profiling for load instructions is performed on all programs. Table 1 shows the statistics from these profiling runs. The number of total profiled load instructions represents the total number of load instructions in each benchmark, as all load instructions are instrumented (profiled). The number of static load instructions represents the number of load instructions that are actually executed. The difference between total profiled and static load instructions is the number of load instructions that are not visited. The number of dynamic load instructions is the total of each ed ic t 0% 10% 20% 30% 40% 50% 70% 80% 90% 100% res s 130 . 132 peg 134 pe rl 147 vo rt e x A ri hm e ic M ean Prediction Accuracies of Load Instructions S tr de Two - Leve Hy b r i d Figure 8. Prediction Accuracy of Load Instructions under Stride, Two-Level, and Hybrid Predictors. load executed multiplied by its execution frequency. Stride, two-level, and hybrid value predictors are simulated during value profiling to evaluate prediction accuracy for each load instruction. Since the goal of this paper is to measure the performance of VSS rather than the required capacities of the hardware buffers, no indices conflicts between loads are modeled. An intelligent index assignment algorithm likely will produce results similar to this, but development of such an algorithm is outside the Total Profiled Load Instructions Static Load Instructions Dynamic Load Instructions 129.compress 96 72 4,070,431 132.ijpeg 5,104 1,543 118,560,271 134.perl 6,029 1,429 4,177,141 147.vortex 16,587 10,395 527,037,054 Table 1. Statistics of Total Profiled, Static and Dynamic Load Instructions. scope of this paper and left for future work. During value profiling, after every execution of a load instruction, the simulated prediction is compared with the actual value to determine prediction accuracy. The value predictor simulators are updated with actual values, as they would be in hardware, to prepare for the prediction of the next use. Each entry for the stride value predictor used has two fields, the stride, the current value. The prediction is always the current value plus the stride. The stride equals the difference between the most recent current values. The stride value predictor always generates a prediction. No finite state machine hardware is required to determine if a prediction should be used. The two-level value predictor design is as in [7], with four data values and six outcome value history patterns in the value history table of the first level. The value history patterns index the pattern history table of the second level. The pattern history table employs four saturating counters, used to select the most likely prediction amongst the four data values. The saturating counters in the pattern history table increment by three, up to twelve, and decrement by one, down to zero. Selecting the data value with the maximum saturating counter value always generates a prediction. The hybrid value predictor of stride and two-level value predictors utilizes the previous description illustrated earlier in Figure 6 of section 2. In the hybrid design, the saturating counters, used to select between stride and two-level prediction, also increment by three, up to twelve, and decrement by one, down to zero. Figure 8 shows the prediction accuracy of load instructions under stride, two-level, and hybrid predictors. The prediction accuracy of the two-level predictor is higher than that of the stride predictor for all benchmarks except 129.compress and 132.ijpeg. However, the average prediction accuracy for the stride predictor is higher than that for the two-level predictor because of the large performance difference in 129.compress. Examining the value trace for 129.compress shows many long stride sequences that are not predicted correctly by the history-based two-level predictor. The hybrid predictor, capable of leveraging the advantages of each prediction method, has the highest prediction accuracy, at 63% on average across all benchmarks. Figures show prediction accuracy distribution for load instructions using the hybrid predictor. Figure 9 is the distribution for static loads and Figure 10 is the distribution for dynamic loads. For 124.m88ksim, 90% of dynamic load instructions have prediction accuracy of 90%. For 129.compress, 80% of dynamic load instructions have prediction accuracy of 90%. For 124.m88ksim, 45% of the static loads have prediction accuracy 90%, representing most of the dynamic load instructions. For 129.compress, 70% of the static loads have prediction accuracy of 90%. These loads are excellent candidates for VSS. Such high prediction accuracy results in low overhead due to the execution of patch-up code. However, for benchmarks 099.go and 132.ijpeg respectively, only 15% and 25% of Figure 9. Prediction Accuracy Distribution for Static Load Instructions Using Hybrid Predictor. red ic tor1030507090-90% -80% -7 0% -60% -5 0% -40% -30% -20% -10% -0% Pred ict ion Accurac ies Percentage of Load 124 . m88ksim 126 . gcc 129 . compress . li 132 . ijpeg 134 . per 147 . vor tex Figure 10. Prediction Accuracy Distribution for Dynamic Load Instructions Using Hybrid Predictor. Hybrid Pred ictor1030507090-90% -80% -70% -60% -50% -40% -30% -20% -10% -0% Prediction Accuracies Percentage of Dynamic Load 126.gcc 129.compress 130. li 132. ijpeg 134.perl 147.vortex dynamic load instructions have prediction accuracy above 50%. Therefore, they will not gain much performance benefit from VSS. The VSS algorithm of Figure 7 is performed on the programs of SPECint95. Prediction accuracy threshold values of 90%, 80%, 70%, 60% and 50% are evaluated. The number of candidates for value-speculative execution is limited to three for each value prediction. This parameter was varied in our evaluation, with the value of three providing good results. For the evaluation of speedup, a very long instruction word architecture machine model based on the Hewlett-Packard Laboratories PlayDoh architecture [13] is chosen. One cycle latencies are assumed for all operations (including LDPRED and UDPRED) except for load (two cycles), floating-point add (two cycles), floating-point subtract (two cycles), floating-point multiply (three cycles) and floating-point divide (three cycles). The LEGO compiler statically schedules the programs of SPECint95. The scheduler uses treegion formation [10] to increase the scheduling scope by including a tree-like structure of basic blocks in a single, non-linear region. The compiler performs control speculation, which allows operations to be scheduled above branches. Universal functional units that execute all operation types are assumed. An eight universal unit (8-U) machine model is used. All functional units are fully pipelined, with an integer latency of 1 cycle and a load latency of 2 cycles. Program execution time is measured by using the schedule length of each region and its execution profile weight. The effects of instruction and data cache are ignored, and perfect branch prediction is assumed in an effort to determine the maximum potential benefits of VSS. Figure 11 shows the execution time speedup of programs scheduled with VSS over without VSS. Five different prediction accuracy thresholds are used to select which load operations are value speculated. The maximum speedup for all benchmarks is 17% for 147.vortex. As illustrated in Figure 10, 147.vortex has many dynamic load operations that are highly predictable. While 147.vortex does not have the highest predictability for load operations, the sheer number, as illustrated in Table 1, results in the best performance. Benchmarks 124.m88ksim and 129.compress also show impressive speedups, 10% and 11.5% respectively, using a threshold of 50%. Speedup for 124.m88ksim actually goes up, even as the prediction threshold goes down, from 90% to 50%. This result can be deduced from the distribution of dynamic loads. For 124.m88ksim, there is a steady increase in the number of dynamic loads available as the threshold decreases from 90% to 50%. There is a tapering off in speedup though, as more miss-predictions are seen near a threshold of 50%. For 129.compress, the step in the distribution of dynamic loads from 80% to 70% is reflected in a corresponding step in speedup. Performance gains for 126.gcc are more reflective of the large number of dynamic load operations than of their predictability. Penalties for miss-prediction at the lower thresholds reduce speedup for 126.gcc. Benchmark 130.li, with a distribution of dynamics loads similar to 126.gcc, has lower performance due to fewer dynamic loads. Benchmark 134.perl clearly suffers 8U Machine Model1.021.061.11.141.18 099.go 124.m88ksim 126.gcc 129.compress 130.li 132.ijpeg 134.perl 147.vortex 90% 80% 70% 60% 50% Figure 11. Execution Time Speedup for VSS over no VSS. Prediction accuracy threshold values of 90%, 80%, 70%, 60% and 50% are used. from not having many dynamic loads. Benchmarks 099.go and 132.ijpeg do not have good predictability for load operations. Based on these performance results, a predictability threshold of 70% appears to be a good selection. From the distribution of predictability for dynamic loads in Figure 10, a threshold 70% includes a large majority of the predictable dynamic loads. Choosing a threshold of predictability lower than 70% results in a tapering off in performance for some benchmarks. This is due to both a higher penalty for miss-prediction and saturation of functional unit resources, resulting in fewer saved execution cycles. 5. CONCLUSIONS AND FUTURE WORK This paper presents value speculation scheduling (VSS), a new technique for exploiting the high predictability of register-writing instructions. This technique leverages advantages of both hardware schemes for value prediction and compiler schemes for exposing ILP. Dynamic value prediction is used to enable aggressive static schedules in which value dependent instructions are speculated. In this way, VSS can be thought of as a static ILP transformation that relies on dynamic value prediction hardware. The results for VSS presents in this paper are impressive, especially when considering that only load operations were considered for value speculation. Future work will include the study of heuristics for selecting register-writing operations in critical paths. Available functional unit resources and remaining data dependencies affect the ability to improve the static schedule and the penalty for patch-up code. VSS should also be applied to operations other than loads based on their predictability and potential benefit to speedup. How many candidates for value- speculative execution (dependent instructions between the predicted instruction and the branch to patch-up code) to allow is also an important parameter. In general, better heuristics for deciding when to speculate values and how many VSS candidates to allow (directly affecting the amount of patch-up code) will be studied. 6. ACKNOWLEDGMENTS This work was funded by grants from Hewlett-Packard, IBM, Intel and the National Science Foundation under MIP-9625007. We would like to thank Bill Havanki, Sumedh Sathaye, Sanjeev Banerjia, and other members in the Tinker group. We also thank the anonymous reviewers for their valuable comments. 7. --R "Dynamic Memory Disambiguation Using the Memory Conflict Buffer," "Value Locality and Load Value Prediction," "Exceeding the Dataflow Limit via Value Prediction," "The Predictability of Data Values," "Value Profiling," "Can Program Profiling Support Value Prediction?," "Highly Accurate Data Value Prediction using Hybrid Predictors," "The Effect of Instruction Fetch Bandwidth on Value Prediction," "Speculative Execution based on Value Prediction," "Treegion Scheduling for Wide-Issue Processors," "The Superblock: An Effective Technique for VLIW and Superscalar Compilation," "Analysis Techniques for Predicated Code," "HPL PlayDoh Architecture Specification: Version 1.0," "Intel, HP Make EPIC Disclosure," --TR The superblock Dynamic memory disambiguation using the memory conflict buffer Value locality and load value prediction Analysis techniques for predicated code Exceeding the dataflow limit via value prediction The predictability of data values Value profiling Can program profiling support value prediction? Highly accurate data value prediction using hybrid predictors Treegion Scheduling for Wide Issue Processors --CTR Dean M. Tullsen , John S. Seng, Storageless value prediction using prior register values, ACM SIGARCH Computer Architecture News, v.27 n.2, p.270-279, May 1999 Tarun Nakra , Rajiv Gupta , Mary Lou Soffa, Value prediction in VLIW machines, ACM SIGARCH Computer Architecture News, v.27 n.2, p.258-269, May 1999 Mikio Takeuchi , Hideaki Komatsu , Toshio Nakatani, A new speculation technique to optimize floating-point performance while preserving bit-by-bit reproducibility, Proceedings of the 17th annual international conference on Supercomputing, June 23-26, 2003, San Francisco, CA, USA Daniel A. Connors , Wen-mei W. Hwu, Compiler-directed dynamic computation reuse: rationale and initial results, Proceedings of the 32nd annual ACM/IEEE international symposium on Microarchitecture, p.158-169, November 16-18, 1999, Haifa, Israel Huiyang Zhou , Jill Flanagan , Thomas M. Conte, Detecting global stride locality in value streams, ACM SIGARCH Computer Architecture News, v.31 n.2, May Compiler controlled value prediction using branch predictor based confidence, Proceedings of the 33rd annual ACM/IEEE international symposium on Microarchitecture, p.327-336, December 2000, Monterey, California, United States Youfeng Wu , Dong-Yuan Chen , Jesse Fang, Better exploration of region-level value locality with integrated computation reuse and value prediction, ACM SIGARCH Computer Architecture News, v.29 n.2, p.98-108, May 2001 Lucian Codrescu , D. Scott Wills , James Meindl, Architecture of the Atlas Chip-Multiprocessor: Dynamically Parallelizing Irregular Applications, IEEE Transactions on Computers, v.50 n.1, p.67-82, January 2001 Martin Burtscher , Amer Diwan , Matthias Hauswirth, Static load classification for improving the value predictability of data-cache misses, ACM SIGPLAN Notices, v.37 n.5, May 2002 Chao-ying Fu , Jill T. Bodine , Thomas M. Conte, Modeling Value Speculation: An Optimal Edge Selection Problem, IEEE Transactions on Computers, v.52 n.3, p.277-292, March

value speculation;instruction level parallelism;VLIW instruction schedulings;value prediction

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An empirical study of decentralized ILP execution models.

Recent fascination for dynamic scheduling as a means for exploiting instruction-level parallelism has introduced significant interest in the scalability aspects of dynamic scheduling hardware. In order to overcome the scalability problems of centralized hardware schedulers, many decentralized execution models are being proposed and investigated recently. The crux of all these models is to split the instruction window across multiple processing elements (PEs) that do independent, scheduling of instructions. The decentralized execution models proposed so far can be grouped under 3 categories, based on the criterion used for assigning an instruction to a particular PE. They are: (i) execution unit dependence based decentralization (EDD), (ii) control dependence based decentralization (CDD), and (iii) data dependence based decentralization (DDD). This paper investigates the performance aspects of these three decentralization approaches. Using a suite of important benchmarks and realistic system parameters, we examine performance differences resulting from the type of partitioning as well as from specific implementation issues such as the type of PE interconnect.We found that with a ring-type PE interconnect, the DDD approach performs the best when the number of PEs is moderate, and that the CDD approach performs best when the number of PEs is large. The currently used approach---EDD---does not perform well for any configuration. With a realistic crossbar, performance does not increase with the number of PEs for any of the partitioning approaches. The results give insight into the best way to use the transistor budget available for implementing the instruction window.

Introduction To extract significant amounts of parallelism from sequential programs, instruction-level parallel (ILP) processors often perform dynamic scheduling. The hardware typically collects decoded instructions in an instruction window, and executes instructions as and when their source operands become available. In going from today's modest issue rates to 12- or 16-way issue, centralized dynamic schedulers face complexity at all phases of out-of-order execution [2] [10]. The hardware needed to forward new results to subsequent instructions and to identify ready-to-execute instructions from the instruction window limits the size of the hardware window. It is very important, therefore, to decentralize the dynamic scheduling hardware. The importance of decentralization is underscored in recently developed processors/execution models such as the MIPS R10000 [20] and R12000, the DEC Alpha 21264 [7], the multiscalar model [4] [14], the superthreading model [17], the trace processing model [13] [15] [19], the MISC (Multiple Instruction Stream Computer) [18], the PEWs (Parallel Execution model [6] [11], and the multicluster model [3]. All of these execution models split the dynamic instruction window across multiple processing elements (PEs) so as to do dynamic scheduling and parallel execution of in- structions. Dynamic scheduling is achieved by letting each execute instructions as and when their operands become available. An important issue pertaining to decentralization is the criterion used for partitioning the instruction stream among the PEs. Three types of decentralization approaches have been proposed based on the criterion they use for this parti- tioning: (i) Execution unit Dependence based Decentralization (EDD), (ii) Control Dependence based Decentralization (CDD), and (iii) Data Dependence based Decentralization (DDD). The first category groups instructions that use the same execution unit-such as an adder or multiplier-into the same PE. Examples are the R10000, R12000, and Alpha 21264. The second category groups control-dependent instructions into the same PE. The multiscalar, superthread- ing, and trace processing models come under this category. The last category groups data-dependent instructions into the same PE. Examples are the MISC, PEWs, and multi-cluster models. Each of the three categories has different hardware requirements and trade-offs. This paper reports the results of a set of experiments that were conducted to provide specific, quantitative evaluations of different trade-offs. We address the following specific questions: What kind of programs benefit from each kind of partitioning ffl How well does performance scale with each decentralization ffl How much benefit would there be if a crossbar is used to interconnect the PEs? The question of how to select the best decentralization approach to use at each granularity of parallelism is an important one, and we discuss how this might be accomplished. Of more immediate concern is the question of whether it is even worth attempting to use such decentralization techniques for more than a few PEs. While we do not yet know the exact shape these execution models will take in the future, we show that if the right choices are made, these decentralization approaches can provide reasonable improvements in instruction completion rate without much of an impact on the cycle time. The rest of this paper is organized as follows. Section 2 provides background and motivation behind decentralization of the dynamic scheduling hardware. It also describes the three decentralization approaches under investigation. Section 3 describes our experimentation methodology. Section presents detailed simulation results of the different decentralization approaches. In particular, it examines the impact of increasing the number of PEs, and the effects of two different PE interconnection topologies. Section 5 presents a discussion of the results, and the conclusions of this paper. Decentralized ILP Execution Models Programs written for current instruction set architectures are generally in control-driven form, i.e., control is assumed to step through instructions in a sequential or- der. Dynamically scheduled ILP processors convert the total ordering implied in the program into a partial ordering determined by dependences on resources, control, and data. This involves identifying instructions that are mutually resource-independent, control-independent, and data- independent. In order to scale up the degree of multiple is- sue, resources that are in high demand are decentralized. To more reordering and parallel execution of instruc- tions, constraints due to resource dependences are overcome (i) by replicating resources such as the fetch unit, the decode unit, the physical registers, and the execution units (EUs) (i.e., functional units), and (ii) by providing multiple banks of resources such as the Dcache, as shown in Figure 1(a). 2.1 Decentralizing the Dynamic Scheduler On inspecting the block diagram of Figure 1(a), we can see that the important structures that remain to be decentralized are the dynamic scheduler (DS), the register rename Dynamic (DS) Scheduler Control Misprediction Information ISA-visible Registers Decode Units Register Rename Unit Control Flow Icache Physical Registers Memory Address Resolution Dcache Banks (a) EU EU EU EU EU EU EU EU Dynamic (DS) Scheduler Dynamic (DS) Scheduler Dynamic (DS) Scheduler Dynamic (DS) Scheduler Control Misprediction Information Flow Control Instruction ISA-visible Registers ICN Distribution Unit Address Dcache Banks Memory Resolution (b) Figure 1: Generic Organization of Dynamically Scheduled ILP Processors (a) Centralized Scheduler; (b) Decentralized Scheduler unit, and the memory address resolution unit 1 . Incidentally, these are the most difficult parts to decentralize because they deal with inter-instruction dependences, which preclude decentralization by mere replication. Of these parts, the DS is the hardest to decentralize because it often needs to handle all of the active instructions that are simultaneously present in the processor. Detailed studies with 0.8-m, 0.35-m, and 0.18-m CMOS technology [10] also confirm that a centralized DS does not scale well. Thus, it is important to decentralize the DS. Many researchers have proposed decentralizing the DS with the use of multiple PEs, each having a set of EUs, as shown in Figure 1(b). In the decentralized processor, the dynamic instruction stream is partitioned across the PEs, which operate in parallel. The 1 The complexity of these structures can be partly reduced by off-loading part of their work to special hardware that is not in the critical path of program execution [8] [9] [19]. instructions assigned to PEs can have both control dependences and data dependences between them. A natural question that arises at this point is: on what basis should instructions be distributed among the decentralized PEs? The criterion used for partitioning the instruction stream is very important, because an improper partitioning could in fact increase inter-PE communication, and degrade performance! True decentralization should not only aim to reduce the demand on each PE, but also aim to minimize the demand on the PE interconnect by localizing a major share of the inter-instruction communication occurring in the processor to within the decentralized PEs. The three current approaches for grouping instructions into PEs revolve around three important constraints to execute instructions in parallel- (i) execution unit dependences, (ii) control dependences, and (iii) data dependences. We shall look at each of the three decentralization approaches. For the ensuing discussion, we use the example control flow graph (CFG) and code shown in Figure 2. This CFG consists of three basic blocks A, B, and C, with block B control-dependent on the conditional branch in A, and block C control-dependent on the conditional branch in B. We shall assume that the control flow predictor has selected blocks A, B, and C to be a trace. I12: BR IF R13 == 0 I9: BR IF R4 >= 0 A I2: I4: BR IF R4 == 0 Figure 2: Example Control Flow Graph and Code 2.2 Execution unit Dependence based Decentralization In this type of decentralization, instructions are assigned to PEs based on the EU that it will execute on. Thus, instructions that are resource dependent on a particular EU execute in the same PE. An artifact of this arrangement is that instructions wait near where its EU dependence is resolved. Interestingly, one of the pioneer dynamic scheduling schemes, implemented in IBM 360/91 [16], had incorporated this type of decentralization in 1967 itself! Very recently the MIPS R10000 and R12000 processors also use this approach [20]. A potential advantage of the EDD approach is that each PE need have only one or a few types of execution units. Another advantage is that instruction partitioning is straight-forward and static in nature when only a single PE has an EU of a particular type. In such a situation, dynamic instances of a given static instruction always get assigned to the same PE. When multiple PEs have an EU of a particular type, then there is a choice involved in allocating instructions that require that EU type. One option in that situation is to do a static allocation by a compiler or by off-line hardware. Another option is to do dynamic allocation (as in Alpha 21264 [7]), perhaps based on the queue lengths in each of the concerned PEs. With either option, a ready instruction may sometimes have to wait for its allotted EU to become free, although another EU of the same type (in another PE) is free. Furthermore, if the processor performs speculative execution, then recovery actions arising from incorrect speculations will necessitate selective discarding of instructions from different PEs. the main shortcoming with the EDD approach, however, is that generally the result from a PE may be needed in any other PE, necessitating a global interconnect between the PEs, which does not scale well [10]. 2.3 Control Dependence based Decentralization (CDD) In the second decentralization approach, a contiguous portion of the dynamic instruction stream is assigned to the same PE. Thus, instructions that are control-dependent on the same conditional branch are generally assigned to the PE to which the branch has been assigned, and instructions wait near where their control dependences will be resolved. Examples for this approach are the multiscalar execution model [4] [14], the superthreading model [17], and the trace processing model [13] [15] [19] 2 . Control-dependence-based decentralization fits well with the control-driven program specification typically adopted in current ISAs. Because control-dependent instructions tend to be grouped together in the program executable, partitioning of instructions among the PEs can be easily done by statically partitioning the CFG. Furthermore, no regrouping of instructions is needed at instruction commit time. CDD hardware implementations proposed so far, such as the multiscalar processor, the superthreading processor, and the trac processors, all organize the PEs as a circular queue as shown in Figure 3. The circular queue imposes a sequential order among the PEs, with the head pointer indicating the oldest active PE. Programs execute on these processors as follows. Each cycle, if the tail PE is idle, the control flow predictor (CFP) predicts the next task in the dynamic instruction stream, and invokes it on the tail PE; a task is a path or subgraph contained in the CFG of the executed program. For instance, if a CDD processor uses trace-based tasks, then blocks A, B, and C of our example code (which forms a trace) are assigned to a single PE. After invocation, the tail pointer is advanced, and the invocation process continues at the new tail in the next cycle. The successor task in our example code will be the one starting at the predicted target of the conditional branch in block C. Thus, the CFP steps through the CFG, distributing tasks (speculatively) Multiprocessors also partition instructions based on control de- pendences. However, their partitioning granularity is generally much coarser (several hundreds of instructions or more per task). Fur- thermore, multiple tasks in a multiprocessor do not share the same register space. to the PEs. When the head PE completes its task, its instructions are committed, and the head pointer is advanced, causing that PE to become idle. When a task misprediction is detected, all PEs between the incorrect speculation point and the tail PE are discarded in what is known as a squash. Figure 3: Block Diagram of an 8-PE CDD Processor Whereas trace processors consider a trace (a single path consisting of multiple basic blocks) as a task, multiscalar processors consider a subgraph of the control flow graph as a task, thereby embedding alternate flows of control in a task. These processors also differ in terms of how the instructions of a task are fetched. Whereas trace processors fetch all instructions of a task in a single cycle and supply them to a PE, multiscalar processors let all of the active PEs parallelly fetch their instructions, one by one. Architectural support is provided in them to facilitate the hardware in determining data dependences. Studies [5] [19] have shown that in CDD processors, most of the register operands are produced in the same PE or a nearby PE, so that a unidirectional ring-type PE interconnect is quite sufficient. Each PE typically keeps a working copy of the register file, which also helps to maintain precise state at task boundaries. 2.4 Data Dependence based Decentralization (DDD) In the third approach of decentralization, data dependences are used as the basis of partitioning. That is, instructions that are data dependent on an instruction are typically dispatched to the PE to which the producer instruction has been dispatched. Mutually data-independent instructions are most likely dispatched to different PEs. Thus, instructions wait near where their data dependences will be resolved. The MISC (Multiple Instruction Stream Computer) [18], the PEWs execution model [6] [11], the dependency-based model given in [10], and the multicluster model [3] come under this category. As data dependences dictate most of the communication occurring between instructions, the DDD approach attempts to minimize communication across multiple PEs. Because the instructions in a PE are mostly data-dependent, it becomes less important to do run-time scheduling within each PE [10]. However, partitioning of instructions in a DDD processor is generally harder than that in a CDD pro- cessor. This is because programs are generally written in control-driven form, which causes individual strands of data-dependent instructions to be often spread over a large segment of code. Thus, the hardware has to first construct the data flow graph (DFG), and then do the instruction parti- tioning, as shown in Figure 4. Notice that if programs were specified in data-driven form, then data-dependence-based partitioning would have been easier. To reduce the hardware complexity, the DFG corresponding to a path (or trace) can be generated by off-line hardware, and stored in a special i-cache for later re-use. I4: BR IF R4 == 0 I9: BR IF R4 >= 0 I12: BR IF R13 == 0 Figure 4: Register Data Flow Graph (RDFG) of Trace ABC in Figure 2 The DDD hardware implementations proposed so far, such as the PEWs [6] [11], the dependence-based model in [10], and the multicluster [3], differ in terms of how the PEs are interconnected. PEWs uses a unidirectional ring-type con- nection, whereas the MISC and dependence-based model of [10] use a crossbar. When a crossbar is employed, all PEs are of same proximity to each other, and hence the instruction partitioning algorithm becomes straightforward. However, as discussed earlier, a crossbar does not scale well. In the multicluster execution model, the ISA-visible registers are partitioned across the PEs. An instruction is assigned a PE based on its source and destination (ISA-visible) regis- ters. Thus, its partitioning is static in nature. In the PEWs execution model, the partitioning is done dynamically. In order to reduce the burden on the partitioning hardware and the complexity on the instruction pipeline, the DFG corresponding to a path is built by off-line hardware, and stored in a special i-cache [11]. Alternately, architectural support can be provided to permit the compiler to convey the DFG and other relevant information to the hardware. 2.5 Comparison We have seen three approaches for partitioning instructions amongst decentralized processing elements. Table 1 succinctly compares the different attributes and hardware features of the three decentralization approaches. From the implementation point of view, CDD and EDD potentially have an edge, because of the static nature of their partition- ing. CDD implementations have a further advantage due Attribute EDD CDD DDD Basis for partitioning Resource usage Control dependence Data dependence Execution unit types in a PE Only a few EU types All EU types All EU types Logical ordering among PEs Partitioning granularity Instruction Task Instruction Time at which partitioning is done Static/Dynamic Static/Dynamic Static/Dynamic Complexity of dynamic partitioning hardware Moderate Moderate High Table 1: Comparison of Different Decentralization Approaches to partitioning at a higher level. Instead of having a 16- way instruction fetch mechanism that fetches and decodes instructions every cycle from an i-cache or a trace cache, the instruction fetch mechanism (including the i-cache) can be distributed across the PEs, as is done in the multiscalar processor [4] [14]. 3 Experimental Methodology The previous section presented a detailed description and comparison of three decentralization approaches. Next, we present a detailed simulation-based performance evaluation of these three decentralization approaches. 3.1 Experimental Setup The setup consists of 3 execution-driven simulators-based on the MIPS-II ISA-that simulate the 3 decentralization approaches in detail. The simulators do cycle-by-cycle sim- ulation, including execution along mispredicted paths. The simulators are equivalent in every respect except for the instruction partitioning strategy. In particular, the following aspects are common for all of the simulators. Instruction Fetch Mechanism: All execution models use a common control flow predictor to speculate the outcome of multiple branches every cycle. This high-level predictor, an extension of the tree-level predictor given in [1], considers a tree-like subgraph of the dynamic control flow graph as the basis of prediction. A tree of depth 4, having up to 8 paths, is used. The predictor predicts one out of these 8 paths using a 2-level PAg predictor. Each tree-path (or trace) is allowed to have a maximum of 16 instructions. The first level table (Subgraph History Table) of the predictor has 1024 entries, is direct mapped, and uses a pattern size of 6. The second level table (Pattern History Table) entries consist of 3-bit up/down saturating counters. A 128 Kbyte trace cache [12] is used to store recently seen traces. The trace cache is 8-way set-associative, has 1 cycle access time, and a block size of 16 instructions. All traces starting at the same address map to the same set in the trace cache. Every cycle, the fetch mechanism can fetch and dispatch up to 16 instructions. Data Memory System: All execution models use the same memory system, with an L1 data cache and a perfect L2 cache (so as to reduce the memory requirements of the simu- lators). The L1 data cache is 64 Kbytes, 4-way set-associative, 32-way interleaved, non-blocking, 16 byte blocks, and 1 cycle access latency. Memory address disambiguation is performed in a decentralized manner using a structure called arcade [11], which has the provision to execute memory references prior to doing address disambiguation. Instruction Retirement: All of the investigated execution models retire (i.e., commit) instructions in program order, one trace at a time, so as to support precise exceptions. PE Interconnection Topology: Three types of PE inter-connects are modeled in the simulators-a unidirectional ring, a bi-directional ring, and a crossbar. The rings take 1 cycle for each adjacent PE!PE transfer. The crossbar takes log 2 p cycles for all PE!PE transfers, where p is the number of PEs. Parameters for the Study: ffl Maximum Fetch Size (f): instructions. ffl PE issue width (d): the maximum number of instructions executed from a PE per cycle is fixed at 3 (be- cause higher values gave only marginal improvements). Thus, each PE has 3 EUs. ffl PE issue strategy: the default strategy is to use out- of-order execution within each PE. The experiments involve varying 3 parameters: the partitioning strategy, the number of PEs (p), and the PE inter-connect 3.2 Benchmarks and Performance Metrics Table 2 gives the list of SPEC95 integer programs that we use, along with the input files we use. The compress95 Average Path Benchmark Input File Trace Length Prediction gcc stmt.i 13.06 81.78% go 9stone21.in 14.29 70.17% li test.lsp 12.28 91.04% vortex vortex.raw 13.59 94.98% Table 2: Benchmark Statistics program is based on the UNIX compression utility, and performs a compression/decompression sequence on a large buffer of data. The gcc program is a version of the GNU C compiler. It has many short loops, and has poor instruction locality. The go program is based on the internationally ranked Go program, "The Many Faces of Go". The li Benchmark Percentage of Instrs Using an EU type EU!EU communication Program Integer Load/Store Branch Int!Int Int!Load/Store Int!Branch Load/Store!Int gcc 43.2% 36.1% 20.7% 26.2% 32.6% 13.9% 11.2% go 52.4% 32.2% 15.4% 34.9% 28.9% 9.0% 14.5% li 26.6% 48.7% 24.7% 14.3% 37.8% 6.3% 7.6% vortex 28.7% 52.4% 18.9% 13.5% 47.6% 7.5% 8.7% Table 3: Distribution of Instructions based on Execution Unit Used program is a lisp interpreter written in C. The m88ksim program is a simulator for the Motorola 88100 processor, and the vortex program is a single-user object-oriented database program that exercises a system kernel coded in integer C. The programs are compiled for a MIPS R3000-Ultrix platform with a MIPS C (Version 3.0) compiler using the optimization flags specified with the SPEC benchmark suite. The benchmarks are simulated to completion or up to 500 million instructions, depending on whichever occurred first. Table also gives some execution statistics, such as the number of instructions simulated, the average tree-path (trace) length, and the path prediction accuracy. From these statis- tics, we can see that gcc and go have very poor control flow predictability, primarily arising from poor instruction local- ity, which causes too many conflicts in the first level table of the predictor. For measuring performance, execution time is the sole metric that can accurately measure the performance of an integrated software-hardware computer system. Accordingly, our simulation experiments measure the execution time in terms of the number of cycles required to execute a fixed number of instructions. While reporting the results, the execution time is expressed in terms of instructions per cycle (IPC). Notice that the IPC figures include only the committed instructions and do not include nops. We also measure register traffic to get more insight into the behavior of the different decentralization approaches. 3.3 Partitioning Algorithms Simulated EDD: In the EDD system, each PE has execution units (EUs) of a particular type. To decide how many PEs should have EUs of a particular type, we measured the percentage of instructions that use each EU type. Table 3 gives these percentages. Based on the percentage of instructions using a particular EU, we used the following EU assignments. When the system has a single PE, all 3 EUs of that PE can execute any type of instruction. When the system has 2 PEs, the first PE houses 3 Integer/FP EUs, and the second PE houses 3 Load/Store/Branch EUs. When the system has 4 or more PEs, the division of PEs is as in Table 4. PEs having EUs Number of Integer Load/Store Branch FP PEs PEs PEs PEs PEs Table 4: Division of PEs for EDD Scheme of the same kind are placed adjacent to each other. The set of PEs with the Load/Store EUs is placed immediately after the set of PEs with the Integer EUs, because there is significant amount of traffic from integer EUs to Load/Store EUs (c.f. Table 3). The instruction partitioning strategy has a dynamic component in that when an instruction can be assigned to multiple PEs, it is assigned to the candidate PE having the least number of instructions. CDD: For studying the CDD partitioning approach, we connect the PEs in a circular queue-like manner. Two different task sizes, namely 8 and 16, are used. In the first case, called CDD-8, a trace of up to 8 instructions is fetched in a cycle and assigned to the PE at the tail of the PE circular queue. In the second case, called CDD-16, a trace of up to instructions is fetched in a cycle and assigned to the tail PE. DDD: For studying the DDD partitioning approach, we use two different partitioning algorithms. The first algorithm, called (DDD-Multicluster), follows the multicluster approach depicted in [3]. A subset of the ISA-visible registers is assigned to each PE such that each ISA-visible register has the notion of a home-PE. For our studies, the n th PE was considered the home-PE for registers r through r is the number of general-purpose registers and p is the number of PEs. The assignment of instructions to PEs is done as depicted in Table 5. Number of Number of PE to which Source Dest. Instruction Registers Registers is Assigned of dest. register of source register of dest. register st source register If 2 or more source registers & destination register are same, then, home-PE of that register; else, home-PE of dest. register Table 5: Instruction Assignment for DDD-M Partitioning Scheme The second DDD algorithm, called DDD-P (DDD-PEWs), makes better use of data dependence information. It uses off-line hardware to construct the register data flow graph (RDFG) for each trace (tree-path) when the trace is encountered for the first time. Once the RDFG of a trace is data dependence chains (or strands) are identified in the RDFG. Some dependence strands may have communication between them. Once the strands are identified, a relative PE assignment is made for the strands, with a view IPC Number of PEs DDD-M EDD IPC gcc Number of PEs31 go IPC Number of PEs31 li IPC Number of PEs31 IPC Number of PEs31 vortex IPC Number of PEs Figure 5: IPC without Nops for Varying Number of PEs, with Unidirectional Ring PE Interconnect to reduce the communication latency between strands. That is, if there is flow of data from one strand to another, the strands are given a relative PE assignment such that the consumer strand's PE is the one immediately following the producer strand's PE. Strands that do not have data dependences with any other strands of the trace are marked as relocatable. At the time of instruction dispatch, the dispatch unit decides the PE placement for each strand based on its dependences to data coming from outside the trace and the relative PE placement decided statically by the off-line hardware. A 2-cycle penalty (stall) is imposed when a trace is seen for the first time in order to form the RDFG and the relative PE assignments. If the PE assigned to an instruction is full, the instruction is assigned to the closest succeeding PE having an empty slot. Performance Results 4.1 IPC with Unidirectional Ring Our first set of studies focuses on comparing the performance of different partitioning algorithms as the number of PEs is varied, and a unidirectional ring is used to connect the PEs. Figure 5 plots the IPC values obtained with the default parameters (PE scheduler for each benchmark. The values of p that we consider are f1, 2, 4, 8, 12, 16g. Each graph in Figure 5 corresponds to a particular benchmark program, and has 3 plots, one corresponding to each decentralization approach. EDD: First of all, the EDD approach does not perform well at all with a ring-type PE interconnect, as expected. This is because the EDD approach is unable to exploit localities of communication, which is very important when using a ring topology to interconnect the PEs. The performance increases slightly as the number of PEs is increased to 2, but thereafter it is downhill. DDD: The performance of the two DDD partitioning algorithms are quite different. The performance of the DDD-M algorithm is generally poor, and similar to the performance of the EDD algorithm simulated. To get good performance from the DDD-M approach, an optimizing compiler needs to rename register specifiers considering the idiosyncrasies of the DDD-M execution model; otherwise, very little of the data dependence localities are likely to be captured. For the DDD-P approach, performance generally keeps increasing as the number of PEs (p) is increased from 1 to 8. For these values of p, the DDD-P algorithm performs the best among the investigated partitioning algorithms. This is because DDD-P is better able to exploit localities of communication when instructions are spread across a moderate number of PEs. The most striking observation is that the performance of DDD-P starts dropping when the number of PEs is increased beyond 8. This drop in performance is because some data-dependent instructions are getting allocated to distant PEs, resulting in large delays in forwarding register values between these distant PEs. One reason for the spreading of data-dependent instructions is that the RDFG formation and the instruction partitioning are done on an individual trace basis. If a knowledge of the subsequent traces is available and made use of while partitioning instructions, then a better placement of instructions can be made. CDD: The performance of both CDD schemes keeps increasing steadily as the number of PEs is increased from 1 to 16. This is because of two reasons: (i) available parallelism increases with instruction window size, and (ii) most register instances have a short lifetime [5] [19], resulting in very little communication of register values between non-adjacent PEs. As the number of PEs is increased beyond 8, the CDD approach starts performing better than the DDD approaches; both DDD and EDD begin to perform worse in this arena! Notice, however, that for three of the benchmarks (compress95, li, and m88ksim), the performance of DDD-P with 4 PEs is better than the performance of CDD- with And for the remaining three benchmarks, the performance of a 4 PE DDD-P processor is not much lower than that of a 16 PE CDD-16 processor. This highlights the importance of developing DDD algorithms that can perform better distribution of instructions over a large number of PEs. 4.2 IPC with Bi-directional Ring To investigate if the unidirectional nature of the ring was the cause of the drop in DDD-P's performance for higher values of p, we also experimented with a bi-directional ring. Table 6 tabulates the IPC values obtained for 12 PE DDD- P with the unidirectional PE interconnect and with the bi-directional PE interconnect. (We simulated the bi-directional ring configuration for 12 PEs, because the performance of DDD-P starts dropping at The data in Table 6 indicate that a bi-directional ring does little to improve the performance of DDD-P when (except for m88ksim which registers a modest improvement from 3.13 to 3.46). Benchmark IPC obtained with Program Unidirectional Ring Bi-directional Ring gcc 2.27 2.27 go 1.82 1.82 li 3.15 3.17 vortex 3.59 3.64 Table Unidirectional Ring and Bi-directional Ring PE Interconnects 4.3 IPC with Crossbar The results presented so far were obtained with ring-type interconnections between the PEs. Next, we investigate how the decentralization approaches scale when the PEs are interconnected by a realistic crossbar. Figure 6 plots the IPC values obtained when the PEs are interconnected by a log 2 p-cycle crossbar. A comparison of the data in Figures 5 and 6 show that with a crossbar interconnect, the performance of EDD has improved slightly for some of the benchmarks. For DDD-P, the performance has decreased (compared to that with ring interconnect) for lower values of p, and remains more or less the same as before for higher values of p. For CDD, the performance with a crossbar is consistently lower than the performance with a ring. In fact, contrary to the case with a ring-type interconnect, the performance of CDD-16 with a realistic crossbar decreases as the number of PEs is increased. Overall, the results with a realistic cross-bar show the performance of DDD-P to be slightly better than that of CDD-16 for most benchmarks. 4.4 Register Traffic In order to get a better understanding of the IPC results seen so far, we next analyze the register traffic occurring in the decentralized processors when different partitioning algorithms are used. Figure 7 plots for the distribution of register results based on the number of PEs they had to travel. For each benchmark, distributions are given for the EDD, CDD-16, and DDD-P partitioning algorithms. The curves for EDD indicate a significant amount of register traffic between distant PEs. For both CDD and DDD, the amount of register traffic between PEs steadily decreases as PE distance increases. For DDD-P, the traffic dies down to almost zero as register values travel about 7 PEs, which explains why using a bi-directional ring does not fetch much of performance improvements. However, a noticeable fraction of register values travel up to 5-6 hops, which affects the performance of the DDD-P scheme. One of the reasons for this is that the DDD-P scheme forms the DFGs for each trace independently, and assigns instructions of a trace to the PEs without considering the DFGs of the subsequent traces who need the values produced by this trace. For CDD, register traffic almost dies down to almost zero, as register values travel about 3 PEs. This is because most register instances have a short lifetime [5] [19], which explains why the performance of CDD with a ring-type interconnect continues to increase as the number of PEs is increased. 5 Discussion and Conclusions The central idea behind decentralized execution models is to split the dynamic execution window of instructions amongst smaller, parallel PEs. By keeping each PE relatively small, the circuitry needed to search it when forwarding newly produced values is greatly reduced, thus reducing the impact of dynamic scheduling on clock speed. By allocating dependent instructions to the same PE as much as possible, communication localities can be exploited, thereby minimizing global communication within the processor. We examined three categories of decentralized execution models, based on the type of dependence they use as the basis for instruction partitioning. These categories are (i) Execution unit Dependence based Decentralization (EDD), (ii) Control Dependence based Decentralization (CDD), and (iii) Data Dependence based Decentralization (DDD). The detailed performance results that we obtained, on an ensemble of well-known benchmarks, lead us to two important conclusions. First, the currently used approach- EDD-does not provide good performance even when the instruction window is split across a moderate number of PEs and when a crossbar is used to connect the PEs. Second, when a unidirectional ring is used to interconnect the PEs, the DDD-P approach provides the best IPC values when a IPC Number of PEs 12Number of PEs gcc IPC EDD DDD-P Number of PEs Number of PEs li31 Number of PEs31 Number of PEs vortex Figure Nops with a Realistic (log 2 p cycle) Crossbar PE Interconnect moderate number of PEs is used. This is due to its ability to exploit localities of communication between instructions. When a large number of PEs is used, the performance of DDD-P starts dropping, and the CDD approach begins to perform better. This is because of the inability of the implemented DDD-P algorithm to judiciously partition complex data dependence graphs across a large number of PEs. Nevertheless, the performance of the implemented DDD-P algorithm with 4 PEs is comparable to or better than the performance of the implemented CDD with PEs. Although the results presented in this paper help in understanding the general trends in the performance of different decentralization approaches, the study of this topic is not complete by any means. There are a variety of execution model-specific techniques (both at the ISA-level and at the microarchitectural level) that need to be explored for each of the decentralized execution models before a conclusive verdict can be reached. In addition, it is important to investigate the extent to which factors such as value prediction, instruction replication, and multiple flows of control introduce additional wrinkles to performance. Finally, it would be worthwhile to explore the possibility of a good blending of the CDD and DDD models by using a DDD-P processor (i.e., a cluster of DDD-P PEs) as the basic PE in a CDD processor. Such a processor can attempt to exploit data independences at the lowest level of granularity and control independences at a higher level. Acknowledgements This work was supported by the US National Science Foundation (NSF) through a Research Initiation Award (CCR 9410706), a CAREER Award (MIP 9702569), and a research grant (CCR 9711566). We are indebted to the reviewers for their comments on the paper and to Dave Kaeli for helps in getting the SPEC95 programs compiled for the MIPS-Ultrix platform. --R "Control Flow Prediction with Tree-like Subgraphs for Superscalar Processors," "Understanding Some Simple Processor- Performance Limits," "The Multicluster Architecture: Reducing Cycle Time Through Partitioning," "The Multiscalar Architecture," "Register Traffic Analysis for Streamlining Inter-Operation Communication in Fine-Grain Parallel Processors," "PEWs: A Decentralized Dynamic Scheduler for ILP Processing," "The Alpha 21264: A 500 MHz Out-of-Order Execution Microprocessor," "Exploiting Fine Grained Parallelism Through a Combination of Hardware and Software Techniques," "Complexity-Effective Superscalar Processors," "Complexity- Effective PEWs Microarchitecture," "Trace Cache: a Low Latency Approach to High Bandwidth Instruction Fetching," "Trace Processors," "Mul- tiscalar Processors," "Multiscalar Execution along a Single Flow of Control," "An Efficient Algorithm for Exploiting Multiple Arithmetic Units," "The Superthreaded Archi- tecture: Thread Pipelining with Run-Time Data Dependence Checking and Control Speculation," "MISC: A Multiple Instruction Stream Computer," "Improving Superscalar Instruction Dispatch and Issue by Exploiting Dynamic Code Sequences," "The MIPS R10000 Superscalar Micro- processor," --TR Exploiting fine-grained parallelism through a combination of hardware and software techniques MISC Register traffic analysis for streamlining inter-operation communication in fine-grain parallel processors The multiscalar architecture Multiscalar processors Control flow prediction with tree-like subgraphs for superscalar processors Trace cache Improving superscalar instruction dispatch and issue by exploiting dynamic code sequences Exploiting instruction level parallelism in processors by caching scheduled groups Complexity-effective superscalar processors Trace processors The multicluster architecture Understanding some simple processor-performance limits The MIPS R10000 Superscalar Microprocessor Multiscalar Execution along a Single Flow of Control The Alpha 21264 The Superthreaded Architecture --CTR D. Morano , A. Khalafi , D. R. Kaeli , A. K. Uht, Realizing high IPC through a scalable memory-latency tolerant multipath microarchitecture, ACM SIGARCH Computer Architecture News, v.31 n.1, March Aneesh Aggarwal , Manoj Franklin, Scalability Aspects of Instruction Distribution Algorithms for Clustered Processors, IEEE Transactions on Parallel and Distributed Systems, v.16 n.10, p.944-955, October 2005 Ramadass Nagarajan , Karthikeyan Sankaralingam , Doug Burger , Stephen W. Keckler, A design space evaluation of grid processor architectures, Proceedings of the 34th annual ACM/IEEE international symposium on Microarchitecture, December 01-05, 2001, Austin, Texas Joan-Manuel Parcerisa , Julio Sahuquillo , Antonio Gonzalez , Jose Duato, On-Chip Interconnects and Instruction Steering Schemes for Clustered Microarchitectures, IEEE Transactions on Parallel and Distributed Systems, v.16 n.2, p.130-144, February 2005 Rajeev Balasubramonian, Cluster prefetch: tolerating on-chip wire delays in clustered microarchitectures, Proceedings of the 18th annual international conference on Supercomputing, June 26-July 01, 2004, Malo, France Balasubramonian , Sandhya Dwarkadas , David H. Albonesi, Dynamically managing the communication-parallelism trade-off in future clustered processors, ACM SIGARCH Computer Architecture News, v.31 n.2, May

speculative execution;instruction-level parallelism;hardware window;control dependence;dynamic scheduling;decentralization;execution unit dependence;data dependence

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Metric details for natural-language spatial relations.

Spatial relations often are desired answers that a geographic information system (GIS) should generate in response to a user's query. Current GIS's provide only rudimentary support for processing and interpreting natural-language-like spatial relations, because their models and representations are primarily quantitative, while natural-language spatial relations are usually dominated by qualitative properties. Studies of the use of spatial relations in natural language showed that topology accounts for a significant portion of the geometric properties. This article develops a formal model that captures metric details for the description of natural-language spatial relations. The metric details are expressed as refinements of the categories identified by the 9-intersection, a model for topological spatial relations, and provide a more precise measure than does topology alone as to whether a geometric configuration matches with a spatial term or not. Similarly, these measures help in identifying the spatial term that describes a particular configuration. Two groups of metric details are derived: splitting ratios as the normalized values of lengths and areas of intersections; and closeness measures as the normalized distances between disjoint object parts. The resulting model of topological and metric properties was calibrated for 64 spatial terms in English, providing values for the best fit as well as value ranges for the significant parameters of each term. Three examples demonstrate how the framework and its calibrated values are used to determine the best spatial term for a relationship between two geometric objects.

Figure 1: Geometric interpretations of the 19 line-region relations that can be realized from the 9-intersection (Egenhofer and Herring 1991). The only other topological invariant used here is the concept of the number of component. A component is a separation of any of the nine intersections (Egenhofer and Franzosa 1995). The number of components of an intersection is denoted by #(A. B). For example, for line-region relation LR 14, #(L. R) 2, whereas for LR 10, #(L. R) =1. The 19 line-region relations can be arranged according to their topological neighborhoods (Egenhofer and Mark 1995a) based on the knowledge of the deformations that may change a topological relation by pulling or pushing the line's boundary or interior (Figure 2). The topological neighborhoods establish similarities that were shown to correspond to groupings people frequently make when using a particular natural-language term (Mark and Egenhofer 1994b). For example, the term crosses was found to correspond to the five relations located in the diagonal from the lower left to the upper right of the conceptual neighborhood diagram (LR 8 to LR 14 in Figure 2). Such groupings of the 9-intersection relations in the conceptual neighborhood diagram may serve as a high-level measure to define the meaning of natural-language spatial relations. However, topology per se may be insufficient as the only measure, particularly in border-line cases where small metric changes have a significant influence on topology. Figure 2: The conceptual neighborhood graph of the nineteen line-region relations(Egenhofer and Mark 1995a). The following sections define two metric concepts-splitting and nearness-that apply to topological relations and may enhance each of the nineteen topological relations to distinguish more details. Splitting Splitting determines how a region's interior, boundary, and exterior are divided by a line's interior and boundary, and vice versa. To describe the degree of a splitting, the metric concepts of the length of a line and the area of a region are used. In the context of topological relations between lines and regions, length applies to the line's interior, any non-empty intersection with a line's interior, or their components; and to region boundaries, any non-empty intersection between a region's boundary and a line's exterior, or their components. Area applies to the interior or regions, the intersections between a line's exterior and a region's interior or exterior, and their components. Among the entries of the 9-intersection for a line and a region, there are seven intersection that can be evaluated with a length or an area (Table 1). Only the three intersections between the line's boundary cannot be evaluated with a length or area measure, because these intersections are 0-dimensional (i.e., points). R R R- length(L. R) length(L. R) length(L. R- ) L- area(L- . R) length(L- . R) area(L- . R- ) Table 1: Area and length measures applied to the nine intersections of the line's interior boundary (L), and exterior ( L- ) with the region's interior ( R), boundary (R), and exterior (R- ). To normalize these lengths and areas, each of them is put into perspective with the line and the region: The two area intersections are compared with the area of the region, resulting in two splitting measures. Another ten splitting measures are obtained by comparing the four length intersections with the length of the line, and the length of the region's perimeter. 3.1 Inner Area Splitting Inner area splitting describes how the line's interior divides the region's interior. With this separation a one-dimensional object splits a two-dimensional object into two (or more) parts such that parts of the region's interior are on one side of the line, and others are located on the opposite side of the line (Figure 3). Inner area splitting only applies to a subset of the 19 region-line relations. Those relations for which the line's interior intersects with the region's interior L. R= ), but the line's boundaries is outside of the region's interior (L . R= ), always have a value for inner area splitting. In addition, inner area splitting may apply if the line's interior intersects with the region's boundary and interior ( L. R= and L. and the line's boundary intersects with the region's interior (L . R= ). In such situations it is necessary that there are more components in the interior-interior intersection than there are components of the intersection between the line's boundary and the region's interior, i.e., #(L. R) > #(L . R). Figure 3: Inner area splitting: the line's interior divides the region's interior into parts on two opposite sides (more complex configurations may have multiple separations on either side of the line). A normalized measure of this property is the inner areasplitting ratio (IAS) as the smaller sum of the areas on either side of the line-left and right are chosen arbitrarily and their choice does not influence the measure-over the total area of the region (Eq. 2). The range of IAS is It would reach 0 if the interior-interior intersection between the line and the region was empty, and is 0.5 if the line separates the region's interior into areas that total the same size on the left-hand side and the right-hand size. min(area(leftComponents(L- . R)), area(rightComponents(L- . R))) area(R) 3.2 Outer Area Splitting Outer area splitting occurs if the line's interior interacts with the exterior of the region such that it produces separations of the exterior between the interior of the line and the boundary of the region. This involves a one-dimensional object that splits a two-dimensional object (the region's exterior) into two (or more) two-dimensional parts: (1) parts of the region's exterior that are bounded because they are completely surrounded by the line's interior and the region's boundary, and (2) parts of the region's exterior that are unbounded (Figure 4). Figure 4: Outer area splitting: the line's interior divides the region's exterior into bounded and unbounded areas (more complex configurations may have multiple areas that are bounded by the same line). Outer area splitting requires that the line's interior intersects with the region's exterior and that the line's boundary is located in the region (L . splitting also may apply to configurations for which line interiors intersect with both the region's interior and boundary ( L. whose line boundaries intersect with the region's exterior (L . For these situations, it is necessary that the region's exterior contains more components of the line's interior than of the line's boundary (#(L. R- ) > #(L . R-)). A normalized measure of outer area splitting is the outerareasplitting ratio (OAS) as the ratio of the sum of the region's area and the bounded exterior, which is the part of the exterior that is enclosed by the line's interior and the region's boundary, over the region's area (Eq. 3). It is greater than zero such that the larger the bounded area, the larger the splitting ratio. It would reach 0 if the bounded area was non-existent (i.e., either an empty intersection between the line's interior and the region's exterior, or an insufficient number of components in the intersection between the line's interior and the region's exterior. area(boundedComponents(L- . R- ) area(R) 3.3 Inner Traversal Splitting The region's interior separates the line's interior into inner and outer line segments. This involves a two-dimensional object splitting a one-dimensional object into two one-dimensional parts (or sets of parts): line parts that are inside the closure of the region, and line parts outside of the region Figure 5). Figure 5: Inner traversal splitting: the region's interior divides the line into parts of inner and outer segments (more complex configurations may have multiple inner and outer segments for a line). Inner traversal splitting applies to relations in which the line's interior is located at least partially in the region's interior ( L. R= ). A normalized measure for the traversal is the inner traversal splitting ratio (ITS) between the length of the inner parts of the line and the length of the total line (Eq. 4). Its range is 0 < ITS 1. ITS would be 0 if the interior-interior intersection between the line and the region was empty. The greatest value is reached if the line's interior is completely contained in the region's interior. length(L. R) length(L) 3.4 Entrance Splitting While the inner traversal splitting normalizes the common interiors with respect to the line's length, the entrance splitting compares the length of the common interiors to the length of the region's boundary. It applies under the same conditions as the inner traversal splitting. Its measure, called the entrancesplittingratio (ENS), captures how far the line enters into the region (Eq. 5). All values of the entrance splitting ratio are greater than zero, but no upper bound exists. length(L. R) length(R) 3.5 Outer Traversal Splitting While the inner traversal splitting describes how much of the line is in the region's interior, the outer traversal splitting refers to the part of the line that is in the region's exterior. Outer traversal splitting applies to relations in which the line's interior is located at least partially in the regions' exterior (L. normalized measure for the traversal is the outer traversal splitting ratio (OTS) between the length of the outer parts of the line and the length of the total line (Eq. 6). length(L. R- ) length(L) 3.6 Exit Splitting Analog to the pair of inner traversal splitting and entrance splitting, the outer traversal splitting has a dual, the exit splitting. It captures how far the line exits the region, and applies under the same conditions as the outer traversal splitting. The exit splitting ratio (EXS) normalizes the length of the line's interior that lays in the region's exterior with respect to the length of the region's boundary (Eq. 7). It is greater than 0 and has no upper bound. length(L. R- ) length(R) 3.7 Line Alongness The region's boundary interacts with the line's interior such that it separates the line into two sets of line parts: line segments that are outside of the region's boundary (i.e., either in the region's interior or exterior), and line segments that are contained in the boundary. This separation makes a one-dimensional object splitting another one-dimensional object into two or more one-dimensional parts ( Figure 6). Figure Line alongness: the region's boundary separates the line's interior into parts of outer and inner segments (more complex configurations may have multiple components in the intersection between the region's boundary and the line's interior). In order to consider line alongness, the line's interior must intersect with the region's boundary As the measure for the separation, we introduce the notion and concept line alongness ratio (LA) as the ratio between the length of all line parts contained in the boundary, and the total length of the line (Eq. 8). The range of the line alongness ratio is 0 LA 1. LA is 0 if the line intersects the region's boundary exclusively in 0-dimensional components, and it reaches 1 if L R. length(L. R) length(L) 3.8 Perimeter Alongness The line's interior separates the region's boundary into two sets of objects, one that coincides whit the line's interior, and another that is disjoint from the line's interior. The separation is such that a one-dimensional object splits another one-dimensional object into two (or more) one-dimensional objects. The perimeter alongness can be measured for relations in which the line's interior intersects with the region's boundary ( L. R = ). The perimeter alongness is measured by the ratio between the length of coinciding parts between the line's interior and the region's boundary and the perimeter, called the perimeteralongnessratio (PA) (Eq. 9). The range of the perimeter alongness ratio is 0 PA < 1. PA is 0 if the interior-boundary intersection between the line and the region consists exclusively of disconnected 0-dimensional components. PA would reach the maximum of 1 if cycles were permitted as lines and such a cycle would coincide with the region's boundary. length(L. R) length(R) 3.9 Perimeter Splitting Perimeter splitting occurs if the line splits the region's boundary into two or more parts. This involves two (or more) zero-dimensional or one-dimensional objects-the line's boundary or interior-cutting another one-dimensional object (the region's boundary) (Figure 7). Figure 7: Perimeter splitting: the line separates the region's boundary into segments (more complex configurations may create multiple segments in the region's boundary). Perimeter splitting requires that the line intersects the region's boundary such that the region's boundary is split into at least two components (#(R- L) 2). The perimeter splitting ratio (PS) is the ratio between the longest of these components and the region's perimeter (Eq. 10). Its range is 0 < PS < 1. max(length(components(L- . R))) length(R) 3.10 Length Splitting While the perimeter splitting compares the length of the longest perimeter component with the total length of the perimeter, the length splitting compares it with the length of the line. The metric measure is the line splitting ratio (LS) (Eq. 11), which is great than 0 without an upper bound. max(length(components(L- . R))) length(L) 3.11 Comparison of the Splitting Ratios Each splitting ratio applies to several different topological relations. Figure 8 shows how the criteria for the ten splitting ratios map onto the conceptual neighborhood graph of the line-region relations (Egenhofer et al. 1993). Each constraint covers a contiguous area. IAS ITS and ENS LA and PA OAS OTS and EXS LS and PS Figure 8: The relations that qualify for inner area splitting (IAS), outer area splitting (OAS), inner traversal splitting (ITS), entrance splitting (ENS), outer traversal splitting (OTS), exit splitting (EXS), line alongness (LA), perimeter alongness (PA), line splitting (LS), and perimeter splitting (PS). Black, gray, and white indicate that the metric measure applies always, sometimes, and never, respectively. 4 Closeness Unlike splitting, which requires coincidence and describes how much is in common between two objects, closeness describes how far apart disjoint parts are. The object parts involved are the boundary and the interior of the line, and the boundary of the region. There is no need to consider the region's interior, since it is delineated by its boundary, and therefore no additional information in could be found by considering it in addition to the region's boundary. Closeness involves considerations of distances among points and lines. For the configurations considered, there are four types of closeness measures of interest (the metric axioms for distances apply, i.e., there is a null element, distances are symmetric, and the triangle inequality holds): (1) the distance between a line's boundary and the region's boundary if the line's boundary is located in the exterior of the region; (2) the distance between a line's boundary and the region's boundary if the line's boundary is located in the interior of the region; (3) the distance of the shortest path between a line's interior and the region's boundary if the line's interior is located in the exterior of the region; and (4) the distance of the shortest path between a line's interior and the region's boundary if the line's interior is located in the interior of the region. The closeness measures are not completely orthogonal, since depending on the shape of the line or the region, they may have the same values. For instance, for the configuration in Figure 9a, the distance from the region's boundary to the line's boundary (i.e., its two endpoints as defined in Section 2) is the same as the distance from the region's boundary to the line's interior, since the line's boundary is the line's closest part to the region's boundary; however, in Figure 9b, the same parameters have different values because the line's interior is closer to the region's boundary than the line's boundary. (a) (b) Figure 9: Two configurations with (a) identical and (b) different values for the distance measures from the line's boundary and interior to the region's boundary. Distances are commonly defined between points; however, the closeness measures require distance measures between a point and a line, or between two lines. 1: The distance between a point p and the boundary of a region (R) is defined as the length of the shortest path from p to R (Eq. 12). $/ q R | dist( p,q) < dist( p,r) Therefore, there is no other point on the region's boundary that would be closer to p Definition 2: The distance between the interior of a line ( L) and the boundary of a region (R) is defined as the length of the shortest path from L to R (Eq. 13). Therefore, there are no other parts in the line's interior that would be closer to any point on the region's boundary 4.1 Outer Closeness The outer closeness describes the remoteness of the region's boundary R from p, a boundary point of a line located in the exterior of the region (Figure 10a). Outer closeness only applies to those line-region relations with at least one point of the line's boundary being located in the region's exterior (L . purely quantitative measure for the remoteness would be the distance between the region's boundary and the line's boundary point(s) in the region's exterior Figure 10b). It is the shortest connections between the line's boundary and the region, i.e., there exists no other point in the region's boundary that would be closer to the line's boundary (Eq. 14). Since this measure is only applicable if L . can never be 0. (a) (b) (c) Figure 10: Outer closeness: (a) the line's boundary in the region's exterior, (b) the remoteness measure BE from the region's boundary to the line's boundary, and (c) the region's outer buffer zone as an equi-distant enlargement of the region. While the actual distance between the two boundaries is a precise measure, it varies significantly with the scale of the representation. For instance, a scaling by a factor of 2 would make any two objects be twice as much remote. A variety of dimension-independent measures could be thought of, such as the proportion by which the line would have to be extended, or shrunken, so that its boundary coincides with the region's boundary. We selected two outer closeness measures: (1) the outer line closeness as the ratio between the distance from the line's boundary to the region's boundary, and the line's length (Figure 10b), and the outer area closeness as the ratio between the area made up by an equi-distant enlargement of the region-also known as a buffer zone (Laurini and Thompson 1992)-and the actual area (Figure 10c). We define the outer area closeness measure (OAC) in terms of the area of the region R and the area made up by the buffer zone, denoted by D (R). It is of width BE and extends into the region's exterior (Eq. 15). OAC is greater than 0 with no upper bound, and would be 0 if BE were 0. The normalization area(D (R))/(area (D (R))+area (R)) would produce values between 0 and 1, however, the distribution would be non linear, particularly for area(R) area(D (R) ). area(D (R)) area(R) The outer line closeness measure (OLC) is defined in terms of BE, the distance from the line's boundary to the region's boundary, and the line's length (Eq. 16). Its values are greater than 0 without an upper bound. It would be 0 if BE were 0. length(L) 4.2 Inner Closeness Analogous to the outer closeness, the inner closeness captures the remoteness of the line's boundary, located in the interior of the region (criterion: L . R= ), from the region's boundary (Figure 11a). The mere distance between the boundaries of the region and the line are captured by a quantitative measure BI (Eq. 17). This distance is greater than 0, because the line's boundary must be located in the region's interior. If both boundary points of the line are inside R, then BI is the distance of that boundary point closest to the region's boundary. BI BI DBI (a) (b) Figure 11: Inner closeness: (a) the line's boundary in the region's interior and (b) the region's inner buffer zone as an equi-distant reduction of the region. min(dist ( p,R)) | p (L . R) (17) The innerareacloseness (IAC) is then defined as the ratio between the area made up by an equi-distant reduction of the region and the actual area (Figure 11b). The buffer zone D (R) has BI the width b and is taken from the region's boundary into the region's interior (Eq. 18). Its rage is D (R) area(R) The inner line closeness (ILC) refers to the relative amount the line has to be extended or shortened to coincide with the region's boundary. The increment is normalized with respect to the line's actual length (Eq. 19). BI length(L) 4.3 Outer Nearness The outer nearness describes how far the line's interior is from the region's boundary (Figure 12a). It only applies to one line-region relation, namely the one with the line's boundary and interior completely contained in the region's exterior (L R- ). The quantitative measure for outer nearness is the length of the shortest connection between the line and the region (Figure 12b). It is always greater than zero, because L must be completely contained in R's exterior (Eq. 20). (a) (b) (c) Figure 12: Outer nearness: (a) the line is completely contained in the region's exterior, (b) the remoteness measure IE from the region's boundary to the line's interior, and (c) the region's outer buffer zone as an equi-distant enlargement of the region. The outerareanearness (OAN) is then defined as the ratio between the area made up of an equi-distant reduction of the region of width IE, denoted by D (R) , and the actual area of the II region R (Eq. 21). OAN's values are greater than 0, with no upper bound. OAN would be 0 if IE were 0. area(D (R)) area(R) The outer line nearness (OLN) normalizes the length by which the line would have to be extended or shortened such that its boundary would coincide with the region's boundary, with respect to the length of the initial line (Eq. 22). The values of the outer line nearness are greater than 0 and increase linearly with the length of IE. length(L) 4.4 Inner Nearness Complementary to the outer nearness, the inner nearness describes how far the line's interior, located in the interior of the region (criterion: L R), is from the region's boundary (Figure 13a) This distance is greater than zero, because the line must be completely contained in the region's interior (Eq. 23). I DII (a) (b) Figure 13: Inner nearness: (a) the line completely contained in the region's interior and (b) the region's inner buffer zone as an equi-distance reduction of the region. The inner area nearness (IAN) is then defined as the ratio between the area made up by a buffer zone of width II, denoted by D (R) , that extends from the boundary into the region's interior II Figure 13b). Its range is 0 < IAN <1 (Eq. 24). area(D (R)) area(R) The inner line nearness (ILN) captures by how much the line would have to be extended in order to intersect with the region's boundary. It is measured as the ratio between the distance to the region's boundary and the length of the line (Eq. 25). The values of the inner line nearness must be greater than zero. II length(L) 4.5 Comparison of the Closeness Measures From the criteria for the closeness measures, one can derive which topological relations may be refined by the corresponding measures (Figure 14). Except for the six topological relations in the bottom triangle of the neighborhood graph, all relations have at least one closeness measure. Those six relations without a closeness measure are such that both line boundaries coincide with the region's boundary, therefore, the distances from the line's parts to the boundaries region are all zero and no refinements can be made to these relations. 5 Parsing and Translating a Graphical Relation into a Verbal Expression With the two sets of parameters we can perform a detailed analysis of a simple spatial configuration with a line and a region, capturing the configuration's topology and analyzing it according to its metric properties. This per se would provide the basis for a computational comparison of two or more spatial configurations for similarity (Bruns and Egenhofer 1996). Here we pursue a different path by mapping the parsed configuration onto a natural-language term that would best describe the spatial relation between the two geometric objects. For the time being, any semantic or presentational aspects (Mark et al. 1995) are ignored in this mapping. The mappings from the topological and metric measures onto corresponding natural-language terms are based on results from human-subject experiments (Shariff 1996). A total of sixty-four English-language terms were tested, for which subjects sketched a road with respect to a given outline of a park such that the sketch would match the corresponding natural-language term that describes the spatial relation. By analyzing the sketches' topological relations and their splitting and closeness measures, we obtained the mappings from the geometry of a configuration onto the corresponding, significant parameters and their values. Significant parameters were distinguished from non-significant ones through a cluster analysis (Shariff 1996). The criterion for a parameter to be considered significant for a specific spatial term was that its standard score was greater than one (i.e., the mean of such a parameter is at least one standard deviation higher than the mean of the entire data set). To demonstrate how the model developed here can be used for such translations, we give three examples in which the spatial relation of a geometric configuration is translated into a natural-language spatial term. 5.1 Example 1 Figure 15 shows a configuration in which a line (e.g., a road) crosses the boundary of a region (e.g., a park). Based on the topology (LR 18), the applicable metric parameters for splitting and closeness are found in Figures 8 and 14, respectively. IAC and ILC IAN and ILN OAC and OLC OAN and OLN Figure 14: The relations that qualify for inner area closeness (IAC), inner line closeness (ILC), outer area closeness (OAC), outer line closeness (OLC), inner area nearness (IAN), inner line nearness (ILN), outer area nearness (OAN), and outer line nearness (OLN). The human-subject tests found that only a subset of these parameters-inner traversal splitting, outer traversal splitting, inner area closeness, and outer area closeness-are significant for the terms that are represented by LR 18. Table 2 shows a sample of eight terms-ends at, ends in, ends just inside, ends outside, enters, goes into, goes out, and goes to-that apply to LR 18, together with the significant parameters. For each parameter, the mean value (i.e., the best fit) and the range of values is given. The value range of a metric parameter refers to the minimum and maximum value obtained from the subjects' sketches. The goal is now to determine which of these terms are a better match for the particular configuration, and which do not convey the meaning the meaning of the configuration. Topological Relation Spatial Term ITS ETS IAC OAC mean range mean range mean range mean range LR ends at 0.24 0.02-0.65 0.76 0.36-0.98 0.79 0.19-0.98 7.69 2.09-20.73 LR ends in 0.51 0.17-0.91 0.49 0.09-0.83 0.81 0.18-0.99 3.66 0.59-9.94 LR just inside 0.16 0.05-0.71 0.84 0.29-0.95 0.55 0.05-0.89 3.95 0.86-10.27 LR LR LR LR LR goes to 0.20 0.04-0.57 0.81 0.44-0.96 0.75 0.18-0.99 7.37 1.63-12.04 Table 2: Spatial terms of topological relation LR 18, with means and value ranges of their significant parameters for splitting and closeness measures. Table 3 summarizes for the four parameters how they are calculated and provides the values obtained for the configuration in Figure 15. inner traversal splitting length(L) I R outer traversal splitting outer area closeness length(L) LI R area(D area(R) R inner area area(D ) closeness area(R) BI R BI Table 3: Calculating the inner traversal splitting, the outer traversal splitting, the outer area closeness, and the inner area closeness for the configuration displayed in Figure 15. By comparing these values with the calibrated model, the terms are ranked according to best fit. The terms ends in, ends outside, enters, goes into, and goes out fall outside of the value ranges of at least two parameters (Table 2) and, therefore, these terms are not considered for this configuration. Among the remaining three terms, ends just inside is the best fit for three parameters; goes to is the second best for three parameters, and ends at ranks third in three out of four times. Therefore, the sentence, The road ends just inside the park would be selected as the best fit, while valid alternatives would be, The road goes to the park or The road ends at the park. R Figure 15: Does the line enter or end just inside the region? 5.2 Example 2 Figure shows a configuration in which a line intersects a region such that it is close to the region's boundary from the inside and farther from the region's boundary in the exterior. A sample of terms that may fit this description are crosses, cuts through, goes through, runs into, and splits. R Figure Does the line cross or cut through the region? For the configuration's topological relation, LR 14, the human-subject tests found two metric parameters to be significant: inner area splitting and outer area closeness. Table 4 displays the mean and the value range for each parameter. Topological Spatial Term IAS OAC Relation mean range mean range LR 14 cuts through 0.32 0.01-0.50 1.75 0.41-6.04 LR 14 runs into 0.13 0.09-0.44 3.42 0.54-11.96 Table 4: Spatial terms of topological relation LR 14, with means and value ranges of their significant parameters for splitting and closeness measures. For the configuration in Figure 16, the term runs into does not qualify, because the configuration is not located within the range of the inner area splitting. From among the remaining four spatial terms, splits comes closest to the mean values of inner area splitting and outer area closeness; therefore, it is selected as the term to describe the configuration. The ranking of the terms in between is more difficult, because they are subject to more subtle differences. Certainly, crosses would be better to describe the scene than cuts through, since both parameters have values that are closer to the mean of crosses than to the mean of cuts through. The term goes through, however, has a better match with the inner area closeness than both crosses and goes through have, however, it ranks considerably lower in the outer area closeness. inner area splitting R 2outer area closeness area(D area(R) R Table 5: Calculating the inner area splitting and the outer area closeness for the configuration displayed in Figure 16. 5.3 Example 3 The following characteristics describe the configuration in Figure 17, in which a line is outside of the region, but follows the shape of the region. Candidate terms to describe this configuration are bypasses, goes up to, and runs along (Table 6). R Figure 17: Does the line run along or bypass the region? Topological Spatial Term OAN OAC Relation mean range mean range goes up to 0.33 0.03-0.75 0.39 0.03-4.79 runs along 0.33 0.16-1.29 1.06 0.30-7.13 Table Spatial terms of topological relation LR 1, with means and value ranges of their significant parameters for splitting and closeness measures. Based on the topological relation, LR 1, the significant parameters are outer area closeness and outer area nearness. The term bypasses does not fall within the ranges of outer area nearness or outer area closeness (Table 7), and is therefore not considered. Both terms goes up to and runs along have the same values for outer area nearness, but since runs along has a significantly lower value for the outer area closeness, it is chosen as the better term to describe the configuration than goes up to. outer area closeness area(D area(R) R outer area nearness area(D area(R) IE R Table 7: Calculating the outer area closeness and the outer area nearness for the configuration displayed in Figure 17. 6 Conclusions This paper developed a computational model to describe the semantics of natural-language spatial terms based on their geometry. The model is based on the 9-intersection topological model and refines is with metric details in the form of splitting and closeness ratios. Splitting ratios describe the proportion of an intersection with respect to the interior or boundary of the two objects. Their normalized values all fall within the interval between 0 and 1 and grow linearly with the size of the intersection. Closeness ratios specify distances between boundaries and interiors. For inclusion or containment relations, the (inner) closeness ratios are normalized to range between 0 and 1, while closeness ratios for disjoint relations are greater than zero with no upper limit. While this may appear to be an inconsistency in the model, it is necessary to obtain measures that grow linearly with the distance between the parts. The model was only developed for relations between a region and a line, however, the concepts generalize to relations between other geometric types, such as two regions or two lines. Splitting and closeness measures can be implemented with standard GIS software. A prototype implementation with the Arc/Info GIS, however, requires the separation of the two objects into different layers (Shariff 1996). A method for computing the intersections necessary to determine the topological relation, using the described by Mark and Xia (1994). In order to determine the metric parameters, AMLs were written to compute intersections, lengths, and areas. Although this method demonstrated the feasibility of implementing the required operators with a commercial GIS, it was cumbersome, because Arc/Info does not support an object concept, and performance was slow. The use of GIS data structures that support an object model, and the integration of algorithms that are tailored to the operations necessary for efficient implementations of the 9-intersection and the metric refinements, are subjects for future investigations. The model developed applies to a number of applications in the area of spatial reasoning, such as similarity retrieval and intelligent spatial query languages. We demonstrated how to use the model to generate natural-languages terms for simple spatial configurations. Based on a calibration of the 9-intersection with splitting ratios and closeness ratios, using human-subjects experiments for sixty-four English-language (Shariff 1996), we showed how a geometric configuration with a linear and an areal object can be analyzed to determine the pertinent features of their spatial relations. Values obtained from this method lead to the selection of appropriate natural-language spatial terms for such spatial scenes. While the splitting and closeness ratios as refinements of topology cover much of the critical properties of the spatial relations, there are other parameters left that may make additional contributions to better choices of natural-language terms. Further investigations-both formalizations and human-subject tests-are necessary to develop a comprehensive and robust set of definitions of the semantics of natural-languages spatial relations. Some of these considerations were outlined in a larger-scale research plan (Mark et al., 1995). The most obvious aspect to study is the influence of the meanings of the objects on the choice of the spatial terms. Whether the objects or concern are roads and parks vs. hurricanes and islands, may lead to different mappings from topology and metric refinements onto the same spatial terms. With respect to geometry, the current model abstracts away all influences of orientation. This is a valid approach for modeling all those concepts and terms that are independent of orientation (such as those based primarily on containment, neighborhood, and closeness; however, orientation, is another parameter that may be critical for those relations expressing information about direction. For example, orientation may be important to distinguish north from south (or above from underneath) Orientations are invariant under translations and scaling, but they may change under rotation. The orientation of the objects can be assessed in several different ways: (1) the global cardinal relation between two objects, i.e., a relation with respect to a fixed orientation framework; (2) the orientation of an individual object, i.e., the cardinal relation between the object's major axis and a global reference frame; and (3) a local relation, i.e., the cardinal direction with respect to the framework established by one of the two objects' orientations. Similar to the metric properties, one could consider purely quantitative measures, e.g., in the form of degrees. Since people usually do not make such a fine distinction, coarser, qualitative models are necessary to formalize the properties of the three orientation concepts. --R Similarity of Spatial Scenes. Cognitive Distance in Intraurban Space. Problems in Cognitive Distance: Implications for Cognitive Mapping. Symbolic Projection for Image Information Retrieval and Spatial Reasoning. A Calculus of Individuals Based on Representing and Acquiring Geographic Knowledge. On the Equivalence of Topological Relations. Categorizing Binary Topological Relationships Between Regions Modeling Conceptual Neighbourhoods of Topological Line-Region Relations Naive Geography. A Critical Comparison of the 4-Intersection and 9-Intersection Models for Spatial Relations: Formal Analysis Qualitative Spatial Reasoning about Distances and Directions in Geographic Space. Using Orientation Information for Qualitative Spatial Reasoning. Language and Spatial Cognition-An Interdisciplinary Study of the Prepositions in English VIsual TRAnslator: Linking Perceptions and Natural Language Descriptions. On the Robustness of Qualitative Distance- and Direction Reasoning A System for Translating Locative Prepositions from English into French. Modeling Spatial Knowledge. Fundamentals of Spatial Information Systems. The Image of a City. Evaluating and Refining Computational Models of Spatial Relations Through Cross-Linguistic Human-Subject Testing Calibrating the Meanings of Spatial Predicates from Natural Language: Line-Region Relations Modeling Spatial Relations Between Lines and Regions: Combining Formal Mathematical Models and Human Subjects Testing. Research Initiative 13 Report on the Specialist Meeting: User Interfaces for Geographic Information Systems. Interaction with Geographic Information: A Commentary. Determining Spatial Relations Between Lines and Regions in Arc/Info Using the 9-Intersection Model Mental Representations of Spatial and Nonspatial Relations. Human Factors in GeographicalInformation Systems. The Measurement of Cognitive Distance: Methods and Construct Validity. The Geometry of Environmental Knowledge. Cognitive Aspects of Human-Computer Interaction for Geographic Information Systems: An Introduction An Algorithm to Determine the Directional Relationship Between Arbitrarily-Shaped Polygons in the Plane The Child's Conception of Space. A Spatial Logic Based on Regions and Connection. A Model of the Human Capacity for Categorizing Spatial Relations. On the Internal Structure of Perceptual and Semantic Categories. Principles of Categorization. Natural Language Spatial Relations: Metric Refinements of Topological Properties. Parts: A Study in Ontology. Algebraic Topology. How Language Structures Space. --TR Representing and acquiring geographic knowledge An algorithm to determine the directional relationship between arbitrarily-shaped polygons in the plane Qualitative Representation of Spatial Knowledge Language and Spatial Cognition Cognitive Aspects of Human-Computer Interaction for Geographic Information Systems Human Factors in Geographical Information Systems The Geometry of Environmental Knowledge Using Orientation Information for Qualitative Spatial Reasoning Natural-language spatial relations --CTR Josef Benedikt , Sebastian Reinberg , Leopold Riedl, A GIS application to enhance cell-based information modeling, Information SciencesInformatics and Computer Science: An International Journal, v.142 n.1, p.151-160, May 2002 Tiago M. Delboni , Karla A. V. Borges , Alberto H. F. Laender, Geographic web search based on positioning expressions, Proceedings of the 2005 workshop on Geographic information retrieval, November 04-04, 2005, Bremen, Germany Haowen Yan , Yandong Chu , Zhilin Li , Renzhong Guo, A Quantitative Description Model for Direction Relations Based on Direction Groups, Geoinformatica, v.10 n.2, p.177-196, June 2006 Salvatore Rinzivillo , Franco Turini, Knowledge discovery from spatial transactions, Journal of Intelligent Information Systems, v.28 n.1, p.1-22, February 2007 Guoray Cai, Contextualization of Geospatial Database Semantics for Human---GIS Interaction, Geoinformatica, v.11 n.2, p.217-237, June 2007

geographic information systems;topological relations;spatial relations;metric refinements;GIS

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Performance Analysis of Stochastic Timed Petri Nets Using Linear Programming Approach.

AbstractStochastic timed Petri nets are a useful tool in performance analysis of concurrent systems such as parallel computers, communication networks, and flexible manufacturing systems. In general, performance measures of stochastic timed Petri nets are difficult to obtain for practical problems due to their sizes. In this paper, we provide a method to compute efficiently upper and lower bounds for the throughputs and mean token numbers for a large class of stochastic timed Petri nets. Our approach is based on uniformization technique and linear programming.

Introduction Stochastic Timed Petri Nets (STPN) are Petri nets where transitions have -ring delays. Since the last decade, they have been receiving increasing interest in the modeling and performance analysis of discrete event systems. Such a tool is particularly useful for modeling systems which exhibit concurrent, asynchronous or nondeterministic behaviors, such as parallel and distributed systems, communication networks and AEexible manufacturing systems. The reader is referred to the extensive survey of [36] on theoretical analyses and applications of Petri nets. Applications to the performance evaluation of parallel and distributed machines (hardware components) and parallel and distributed computations (software components) can also be found in [3] and the special issue of J. of Parallel and Distributed Computing (Vol. 15, No. 3, July 1992). Most literature of STPN is on Stochastic Petri Nets (SPN) [29, 35], where transition -ring times are mutually independent exponentially distributed random variables, and their extensions: Generalized Stochastic Petri Nets (GSPN) [2] where immediate transitions (i.e. those without -ring delay) are allowed, and Extended Stochastic Petri Nets (ESPN) [28] where transitions are allowed to generate random numbers of tokens upon -rings. Numerical analysis of such nets is based on the analysis of the embedded Markov chains. Decomposition techniques are proposed, see e.g. [19, 34] and references therein. Analytical solutions exist in product-form for equilibrium distributions for special cases of SPN, see [15] and references their in. There also exist analyses of stochastic timed Petri nets without Markovian assump- tions. Most of them provide performance bounds, see [10, 11, 17, 18, 25]. Others analyze stability conditions [4, 9]. The reader is referred to [5] for a survey on recent results on quantitative analysis of STPN, including approximations and simulations. Although there exist various quantitative analysis techniques and some software tools (e.g. GreatSPN [23] and SPNP [27]) for STPN, the applications of STPN are most often limited to small size problems. This is mostly due to the time and space complexity of numerical analysis algorithms and of simulations. In this paper, we provide a new method to compute eOEciently upper and lower bounds for linear functions of the throughputs and mean token numbers in general Markovian Petri nets. Our approach is based on uniformization technique and linear programming. The STPN models under consideration are closely related to GSPN models de-ned in [24], with in addition the possibility of randomly generating tokens upon transition -rings. Uniformization technique is one of the most useful techniques for analyzing continuous time Markov chains [31]. In [32], such a technique was used to establish linear equality constraints among the expectation of state variables in queueing networks. This allowed RR n\Sigma2642 4 Z. Liu the authors to bound the performance measures, both above and below, by solving a linear program. Similar approaches were taken to determine lower bounds on achievable performance of control policies in multiclass queueing networks [13], optimal control policies for Klimov's problem [14], and stability regions of queueing networks and scheduling policies [33]. In these studies, linear or nonlinear programming were used to obtain bounds. The method of linear programming has already been used in operational analysis for deriving bounds in non-Markovian STPN [17, 18, 25]. Since no statistical assumptions are made on the distributions of -ring times, such bounds are usually loose. Several techniques were proposed for the improvement of such bounds in special cases of Petri nets [20, 21]. In our work, we consider Markovian STPN. We show that, like in [32, 13], the Markovian assumption allows us to establish a set of linear equality constraints among the expectation of state variables in the Petri nets, such as token numbers in the places and indicator functions of whether transitions are enabled. More precisely, we analyze the evolution of state variables in steady state and write out evolution equations using the uniformization technique. Taking the quadratic forms of these equations allows us to establish the linear constraints. Exploiting further structural and probabilistic properties of the Petri nets, we obtain an augmenting set of linear equalities and inequalities, some of which are similar to those in [25]. Upper and lower bounds of performance measures are then obtained by solving the linear program. The paper is organized as follows. In Section 2, we de-ne the STPN models under consideration as well as the notation. In Section 3, we derive the linear equalities based on the uniformization technique. In Section 4, we establish other linear constraints based on the behavioral properties and probabilistic laws. In Section 5, we provide the summary of the linear programming formulation. In Section 6, we present applications of our technique. Finally, in Section 7, we conclude with remarks on the extensions of our results. Notation A Petri Net can be viewed as a directed graph E), where the set of vertices is the union of the set of places P and the set of transitions T . The set of arcs E is composed of two subsets E 0 and E 00 . The arcs of E 0 are either of the form (p; t) or of the form (t; p) with We shall denote by the set of transitions that precede place p in P: ffl the set of transitions that follow place p in P: the set of places that precede transition t in T INRIA Performance Analysis of Stochastic Timed Petri Nets 5 the set of places that follow transition t in g. The arcs of E 00 are inhibitor arcs connecting places to transitions. For any t 2 T , let ffi t be the set of places from which there is an inhibitor arc, and for any p 2 P, let p ffi be the set of transitions to which there is an inhibitor arc. Denote by j p;t the weight of the inhibitor arc from place p to transition t, t. The net N is strongly connected if there is a path from any place/transition to any place/transition. For all p 2 P, de-ne the following set of transitions: Tokens circulate in the Petri Net. This circulation takes place when transitions are -red. When transition t 2 T is -red, - p;t tokens are consumed at each place p 2 ffl t, and oe t;p tokens are created at each place p 2 t ffl . Variables - p;t and oe t;p are considered as the weights of the arcs of E 0 . An example of the Petri net is illustrated in Figure 1. It contains 7 places g. Transitions t are immediate transitions. Places p 1 and p 6 have initial marking 1, whereas the others have initial marking 0. There are two inhibitor arcs (p represented by arcs ended with a circle. When the weights of the arcs are upper bounded by 1, N is called an ordinary net, as opposed to weighted net. In this paper, we will consider a more general case where the numbers of tokens created by -ring completions are random variables. When transition t 2 T is -red for the n-th time, oe t;p (n) tokens are created at each place p 2 t ffl . For all is assumed to be a sequence of independent and identically distributed (i.i.d.) random variables. The sequences of random variables foe t 1 ;p (n); are, however, in general dependent for t 1 6= t 2 . Let oe t;p be the expectation of oe t;p (n). For all (n) and oe t;p 2 (n) can be dependent if p 1 . For example, when creates one token in one of its output places after each RR n\Sigma2642 6 Z. Liu Figure 1: An example of Petri net. -ring. Two cases will be considered: independent token generation and selective token generation. In the case of independent token generation, we assume that for any t 2 T , the sequences of random variables foe t;p (n)g n , are assumed to be (statistically) independent. In the case of selective token generation, however, the sequences of random variables foe t;p (n)g n , are dependent in such a way that for all n, at most one of the output places has tokens created: so that oe t;p 1 for any p 1 special case of selective token generation is the routing mechanism where a token is generated at one and only one of the output places after each -ring: below discussions on immediate transitions). There are two special classes of ordinary Petri nets, referred to as state machines and marked graphs. A state machine is an ordinary Petri net without inhibitor arcs such that for each transition t, ffl t is a singleton and marked graph is an ordinary Petri net without inhibitor arcs such that for each place p, both ffl p and p ffl are singleton. Firings of transitions are timed, i.e., each -ring takes a certain amount of time before completion. The token consumptions in places of ffl t and token creations in places of t ffl occur simultaneously at the end of a -ring of transition t. Throughout the paper we will assume that all the -ring times are independent random variables. The -ring times of transition are i.i.d. random variables of exponential distribution with parameter - t . In GSPN framework, Petri nets can have immediate transitions, i.e. transitions whose -ring times are zero. In this case, immediate transitions have higher -ring priorities, see [24]. Using algorithms of [26], these immediate transitions can be eliminated without changing performance behavior of the net. INRIA Performance Analysis of Stochastic Timed Petri Nets 7 Of particular interests are immediate transitions which play roles of synchronization and/or routing. More precisely, in this case, we assume that for any immediate transition t, t is the only output transition of all its input places, i.e. t. Further, we assume that for any immediate transition t, t, and ffl either oe t;p 0 ffl or oe t;p with a harmless abuse of notation, the index ffl p denotes the unique transition preceding place p. In the Appendix, we present a direct transformation technique which removes this kind of immediate transitions without changing the -ring behavior of the other transitions. Thus, we will assume throughout this paper that the Petri net N has no immediate transition, so that all parameters - t are -nite. A transition t is enabled to -re when there are at least - p;t tokens at each place p 2 ffl t and there are at most j tokens at each place t. We adopt the single-server semantics for the transitions. A -ring can start only if the transition is enabled and the previous -ring has completed. It is assumed that -rings are started as soon as possible. The case of in-nite-server semantics will be discussed in Section 7. A -ring of transition t is preempted when the transition is disabled (i.e. at least one place strictly less than - p;t tokens, or at least one place has more than or equal to j p;t tokens) before the -ring time expires. The -ring is resumed as soon as the transition becomes enabled. The disabling of a transition is due both to competitions with other transitions having common input places (some tokens in these places can be consumed by other transitions during the -ring of the transition), and to token arrivals in input places of inhibitor arcs. The -ring mechanism described here is called (cf. [1]) race policy with age memory. Note that for the case of exponential distributions of -ring times, the race policies with or without age memory have stochastically the same performance behavior due to the memorylessness property of exponential distributions. However, in Section 7, when we consider the case where -ring times have general distributions, the race policy under consideration will be that with age memory. The state of the system is characterized by the marking is the number of tokens in place p at time - . The process X(-) is assumed to be left-continuous so that X p (-) is the number of tokens in place just before time - . The initial marking is the marking at time 0. RR n\Sigma2642 8 Z. Liu The Markovian Petri net described above will be denoted by Throughout this paper we will assume that the Petri net is live. Moreover, we assume that the net is stable in the sense that X(-) converges to a stationary variable X (of dimension jPj) when - goes to in-nity. Moreover, we assume that the -rst and second moments of X are -nite, i.e. E[X p assumptions it is easy to see (using H#lder's inequality) that for all 1. Let e t (-) be the indicator function of whether transition t is enabled at time - (or more precisely, just before time - Let e t be the stationary version of e t (-), and q Denote by x the mean number of tokens in place p 2 P, and y . The corresponding vectors are denoted by the (asymptotic) throughput of transition t 2 T , i.e. the number of completed -rings of transition t per unit of time, and In the sequel, we provide a method of computing upper and lower bounds of L(x; for any arbitrarily -xed linear function L. Our approach is based on linear programming. The upper (resp. lower) bound is obtained by maximizing (resp. minimizing) the objective function L under linear constraints. 3 Uniformization and Linear Equalities We will use the uniformization technique to derive linear equalities between variables x, y, q and '. We will consider the Petri net N where each transition t 2 T is continuously -ring with i.i.d. exponentially distributed -ring times of parameter - t . When a -ring is completed at transition t 2 T , there are two possibilities. If t is enabled, then tokens are consumed in places ffl t and are created in places t ffl . Otherwise, if t is disabled when the -ring is completed, nothing happens, and this -ring completion corresponds to a -ctive -ring completion. INRIA Performance Analysis of Stochastic Timed Petri Nets 9 Let f- n g be the sequence of time epochs of, real or -ctive, -ring completions in N . It is clear that f- n g is distributed according to a Poisson process with parameter Let F -n denote the oe-eld generated by the events up to time - n . Let A t (n) be the indicator function such that A t only if the n-th, real or -ctive, -ring completion occurs at transition t 2 T . Clearly, for any t 2 T , fA t (n)g is a sequence of i.i.d. random variables, independent of e t (- n ), such that P Since for any -xed t 2 T , the random variables (oe t;p (n); are i.i.d. in n, we can assume with no loss of generality that the numbers of tokens created in places t ffl at time - n are oe t;p (n), is enabled at time - n . We assume without loss of generality that the system is in steady state so that, owing to process see time average) property (cf. e.g. [6]), (X(- n ); e(- n )) has the same law as (X; e). The throughput of transition t 2 T can be computed as follows. In the system, transitions are -red, either really or -ctively, at the rate of -. At each -ring completion epoch - n , the -ring occurs at transition t 2 T with probability - t =-. Therefore, (real or -ctive) -ring completions occur at transition t at the rate of - t . Since these -ring completions are independent of e t , we have The following evolution equation is essential. For all (2) Taking the conditional expectation yields RR n\Sigma2642 Z. Liu Thus, In the steady state, E[X p (- n+1 so that by taking expectation in (3), we obtain the following AEow balance equalities: Calculating the second moments from (2) yields Thus, In the steady state, E[X 2 and E[e t (- n Hence, by taking expectation in (5), we obtain the following second moment condition:X INRIA Performance Analysis of Stochastic Timed Petri Nets 11 More generally, for any p 1 we compute the expectation of the product of numbers of tokens from (2). Assume -rst that token generations of all transitions t 2 T are statistically independent i.e., random variables oe t;p (n), are independent. Then: RR n\Sigma2642 Z. Liu After some simple algebra, we obtain In the steady state, Thus, by taking expectation in (7) we obtain the following population covariance condition: INRIA Performance Analysis of Stochastic Timed Petri Nets 13 Note that when relations (6) and (8) are identical. Assume now that token generations of some transitions are selective, i.e. for some t 2 be the subset of transitions which have selective token generations. Then, for any t 2 T 0 and any p 1 Therefore, by a similar computation we obtain: Observe that equality (8) can be considered as a special case of (9). Indeed, if T then both equalities coincide. 4 Other Constraints In this section, we derive other linear constraints of variables x, y and q. Except for (24), the linear constraints established in this section requires no Markovian assumption and holds for general stochastic Petri nets. 4.1 Behavioral Properties Liveness. Since we assume that the net is live, we have that for any - , at least one transition is enabled, so that RR n\Sigma2642 14 Z. Liu As a consequence, so that ConAEicting transitions. For all where, by convention, - t, and j t. For any pair of transitions enabled only if t 1 is enabled, so that e t 1 (-) for all - . Hence, If transitions are in equal conAEict, i.e. -(t 1 then the above relation implies that q t 1 Boundedness. For all be the minimum and maximum numbers, respectively, of tokens in place p. Then, trivially, As a consequence, for any place p 2 P such that The bounds (13) can be extended to a set of places S ' P. Let b S - 0 and be the minimum and maximum of total numbers of tokens in places of S. Then, trivially, Cycle population conservation. A special case of (15) is when the subset of places consists of a cycle, i.e., there is a set of transitions INRIA Performance Analysis of Stochastic Timed Petri Nets 15 such that ffl g. Since the net is live and stable, the sum of tokens in these places is constant: Denote by C any cycle in N , and C C the population in C. It then follows Reachable markings. Let be the incidence matrix such that C E). It is well-known (see e.g. [36]) that any reachable marking X from the initial marking M can be written as where the superscript T denotes the transpose operator, and the (column) vector H corresponds to the -ring sequence to reach X (or more precisely, the vector of numbers of -rings of each transition in order to reach X). Let X in (19) be the random variable of the marking in the stationary regime. Then, by taking expectation in (19) we obtain where are newly introduced unknown variables. Rewriting (20) in scalar form yields oe t;p u 4.2 Constraints Derived from Probability Theory path comparisons. Since for any t 2 T , e t - 1 almost surely (a.s.), the enabling rate is bounded by one: For the same reason, we have X p e t - X p a.s., so that RR n\Sigma2642 Z. Liu Another consequence is A t so that According to the relation so that, by taking the expectation, we obtain Probabilistic inequalities. According to Chernooe's inequality, we get for all n - 1, where V is the iminj operation. For bounded places INRIA Performance Analysis of Stochastic Timed Petri Nets 17 Therefore, or, equivalently, Consider any transition t 2 T such that all incoming places are bounded, i.e., for all 1. Using the fact that we obtain that where the last inequality comes from relations (27,29). Hence, we obtain an enabling lower Applying again Chernooe's inequality to (30) yields where the last inequality comes from relations (27,29). Thus, we obtain an enabling upper RR n\Sigma2642 Z. Liu Note that in (32), the iminj operator i is nonlinear. However, linear inequalities can be generated by taking either operand of any of the iminj operators. Consider now an arbitrary bounded place p with bound B p . Then for any t 2 T , Thus, Similarly, where the last inequality comes from (28). Thus, Single entry transitions. For all INRIA Performance Analysis of Stochastic Timed Petri Nets 19 so that Little's Law. According to Little's law (see e.g. [40]), for all is the input rate tokens at place p, and R p is the mean token sojourn time at place p. Since R p is lower bounded by the minimum -ring times of output transitions of p, we obtain or, equivalently, 4.3 Subnet Throughputs Like in [21], we derive bounds on throughputs of transitions by comparing throughputs of N with those in the subnets (when they are considered in isolation) of N . We will consider in particular two special classes of subnets: strongly connected state machines (SCSM) and strongly connected marked graphs (SCMG). E) be an arbitrary Petri net, and N subnet of N , is a restriction of E on fP 0 S T 0 g \Theta fP 0 S T 0 g. Assume that the transitions of T 0 (resp. arcs of E 0 , places of P 0 ) have the same sequences of -ring times (resp. weights, initial markings) in both nets. Assume further that none of the places of connected with transitions of in the original net N by non-inhibitor arcs, i.e. in N , there is no (t; p) 2 t denote the throughput of transition t 2 T 0 when the subnet N 0 is considered in isolation. The following theorems show that under some conditions, the throughputs of these transitions of the N 0 are upper bounds of the throughputs of the same transitions in the original net. Theorem 1 If N 0 is a strongly connected marked graph, then for any transition t in N 0 , t . Proof. Due to the fact that in the original net N , none of the places of P 0 is connected with transitions of non-inhibitor arcs, the subnet is connected with the rest of RR n\Sigma2642 Z. Liu the system only through transitions of T 0 . As N 0 is a strongly connected marked graph, no transitions in N 0 are in conAEict. Moreover, in N , the -ring mechanism is race policy with age memory. Thus, for any transition t 2 T 0 , the only eoeect that tokens in places of have is delaying the -rings of t of T 0 . Thus, by the monotonicity property of marked graphs [11], we conclude that ' t - ' 0 t for all t 2 T 0 . Theorem 2 Assume that N 0 is a strongly connected state machine such that for any two transitions t 1 and t 2 of N 0 , t 1 and t 2 are in conAEict in N 0 implies that t 1 and t 2 are in equal conAEict in N , i.e. ffl t . Then for any transition t in N 0 , ' t - ' 0 t . Proof. The proof is similar to that of Theorem 1. Note -rst that the subnet N 0 is connected with the rest of the system only through transitions of T 0 or inhibitor arcs. Under the assumption of the theorem, any two transitions which are in (equal) conAEict in N 0 are in equal conAEict in N . Moreover, in N , the -ring mechanism is race policy with age memory. Thus, in N , tokens in the places of [ t2T 0 not change the winners of -ring races among transitions of T 0 . In other words, the only eoeect that tokens in places of ffl have is delaying the -rings of transitions of T 0 . Thus, by the monotonicity property of state machines [8], we conclude that ' t - ' 0 t for all t 2 T 0 . Note that the above two theorems hold for any arbitrarily -xed sequences of -ring times. No Markovian assumption is needed. EOEcient computational algorithms for computing the throughput of state machines have been proposed in the queueing literature, see e.g. [16, 22, 38, 39]. The computation of throughput of SCMGs has been investigated in [11], where various computable upper bounds were proposed. Exact value of the throughput can be obtained by simulation using matrix multiplications in the (max; +) algebra, see [7]. 5 Summary of the Linear Programming Formulation Theorem 3 Let be an arbitrary Markovian timed Petri net, and an arbitrary linear function de-ned on the nonnegative state variables x; of the net. Let ff and fi be the solutions of the linear programming problems INRIA Performance Analysis of Stochastic Timed Petri Nets 21 such that the linear constraints of Table (1) are satis-ed, where u Recall that in Table 1, inequalities containing the operator i inequalities generated by taking either operand of any of the iminj operators. 6 Applications In this section we illustrate applications of the above techniques to the performance ana- lyses. We shall consider two applications, one in manufacturing system, another in parallel computing. Unless otherwise stated, the numerical results are obtained without linear inequalities pertaining to boundedness of subsets of places and subnet throughputs. 6.1 Production Line The -rst example is concerned with a production line with in-nite supply, see Figure 2- (a). In the example, there are four servers, represented by circles. The -rst server has an in-nite-capacity buoeer with an in-nite number of production requirements, represented by small dashed circles. The other three servers have -nite-capacity buoeers: 3, 2 and 4. For starts a service only when the downstream buoeer 1 has at least one empty room. This corresponds to the so-called blocking before service. The corresponding Petri net model is depicted in Figure 2-(b), where transitions re-present the servers and the initial markings of the places on the bottom represent the buoeer capacities. We assume that the service times at server i are i.i.d. exponentially distributed with parameter 4. The objective function in this problem is the total throughput . The numerical results are presented and compared with the exact values. In the experimentation, we have carried out computations for -ve sets of parameters of - i 's. The lower and upper bounds are given in the columns il.b.j, iu.b.1j and iu.b.2j, whereas the exact values are provided in the column iexactj. The upper bounds in column iu.b.1j are obtained by further using subnet throughput constraints. In columns io.l.b.j and io.u.b.j, we also present the bounds computed by linear programming approch based on the linear constraints of Section 4 without Markovian assumption (which implies in particular that the linear equalities of Section 3 are not used). RR n\Sigma2642 22 Z. Liu Table 1: Summary of Linear Constraints AEow balance second moment 2 population covariance liveness conAEicting transitions boundedness cycle population reachable marking x sample path comparisons y p;t - x p 8p 2 enabling bound q t Bp \Gamma- p;t +1 xp Bp \Gammaj p;t +1 INRIA Performance Analysis of Stochastic Timed Petri Nets 23 (a) (b) server 3 server 4 server 2 server 1 Figure 2: (a): Example of production line. (b): The corresponding Petri net model. Table 2: Bounds on the throughput of the production line Case exact u.b.1 u.b.2 o.l.b. o.u.b. 5 1.111 1.111 1.111 1.111 1.350 2.667 2.963 2.963 1.111 4.444 RR n\Sigma2642 Z. Liu Recall that the Petri net is a marked graph so that according to [11] the throughput is increasing in the -ring rates of transitions. Such a fact is clearly shown in the column iexactj for cases 1, 2, 3 and 4. It is worthwhile noticing that the lower and upper bounds in the columns il.b.j, iu.b.1j and iu.b.2j also reAEect such monotonicity. 6.2 Cyclic Execution Consider now performance analysis of a parallel computing system. Parallel programs are represented by directed acyclic graphs, referred to as task graphs, where vertices correspond to tasks of a parallel program, and directed edges correspond to precedence relations between tasks: a task can start execution only when all its predecessors have completed execution. The tasks are assigned to the parallel processors for execution according to some prede-ned rules. In our example, parallel programs have the same structure, given by the task graph in Figure 3-(a). These programs dioeer only in the running times of tasks which are independently and exponentially distributed random variables, with parameters for tasks There are three identical processors. Tasks 1 and 2 are assigned to processor 1, tasks 3 and 4 to processor 2, and tasks 5 and 6 to processor 3. (a) (b) Figure 3: (a) Task graph of parallel programs. (b) Cyclic execution of the parallel program. On each processor, dioeerent instantiations of the same task are executed according to the rule -rst come -rst serve (FCFS). i.e., task i of the n-th arrived program can start execution only after task i of program completes. Dioeerent tasks assigned to the same processor are, however, executed according to the processor sharing (PS) discipline. In our example, since only two dioeerent tasks are assigned to each processor, the processor INRIA Performance Analysis of Stochastic Timed Petri Nets 25 is shared by at most two tasks. A parallel program is considered completed if all its tasks -nish their execution. We consider cyclic execution of the task graph, cf. Figure 3-(b). The cyclic execution is de-ned in such a way that task t1 (resp. task t2) of program n + h can start execution only after task t5 (resp. t6) of program n completes execution, The number h is referred to as the height of the cyclic execution in the literature [30]. The representation of this parallel computing system by STPN is illustrated in Figure 4. The initial marking of place p1 (the same for place p2) corresponds to the height h. p3 p6 p7 p8 Figure 4: Petri net representation of the cyclic execution Inhibitor arcs are used to model the PS mechanism. Transitions t11 and t12 (resp. t21 and t22, t31 and t32 and t33, t41 and t42, t51 and t52, t61 and t62 and t63) represent tasks 1 (resp. 2, 3, 4, 5, 6). Firing times of transitions t11 and t12 (resp. t21 and t22, RR n\Sigma2642 26 Z. Liu t31 and t32 and t33, t41 and t42, t51 and t52, t61 and t62 and t63, are exponentially distributed with parameter Two or three transitions are used for each task in order to represent situations whether the execution of a task is shared with others. Note that transitions t32 and t33 (resp. t62 and t63) are never enabled simultaneously. The use of two additional transitions for task 3 (resp. task is due to the fact that in each program, task 4 (resp. task 5) is allowed to start execution only when both tasks 1 and 2 (resp. tasks 3 and 4) have completed execution. Thus, the execution of task 3 (resp. task 6) is not shared with task 4 (resp. task 5) if only task 1 or only task 2 (resp. only task 3 or only task The objective function in this problem is still the total throughput, i.e. the sum of transition throughputs. It is easy to see that this total throughput is equal to six times the throughput of the parallel system in terms of the number of programs completed per unit of time. In Table 3, we provide numerical results for -ve dioeerent sets of parameters with -xed height 4. The lower and upper bounds are given in the columns il.b.j and iu.b.j, whereas in columns io.l.b.j and io.u.b.j, we also present the bounds computed by linear programming approch based on the linear constraints of Section 4 without Markovian assumption (which implies in particular that the linear equalities of Section 3 are not used). Table 3: Bounds on the throughput of the cyclic execution Case In Figures 5 and 6, we provide the curves of the bounds as functions of the the height of the cyclic execution in Cases 1 and 2. 7 Conclusions and Extensions In this paper, we have established performance bounds for Markovian STPN by taking a linear programming approach. We -rst provided a set of linear equality constraints among the expectation of state variables in the Petri nets, such as token numbers in the places and indicator functions of transition enabling. We further obtained an augmenting set of INRIA Performance Analysis of Stochastic Timed Petri Nets 27 O.L.B. L.B. U.B. O.U.B. height (h) Figure 5: Bounds as functions of height of the cyclic execution for Case 1 O.L.B. L.B. U.B. O.U.B. height (h) Figure Bounds as functions of height of the cyclic execution for Case 2 RR n\Sigma2642 28 Z. Liu linear equalities and inequalities by exploiting structural and probabilistic properties of the Petri nets. These linear constraints allowed us to compute upper and lower bounds of performance measures by solving the linear program. We have applied this method to performance analyses of a manufacturing system and a parallel system. The constraints derived in Section 4 are not restricted to exponentially distributed -ring times. These inequalities can also be used in operational analysis of timed Petri nets. In Theorems 1 and 2, we compared throughputs of transitions in a net and those in a SCMG or a SCSM subnet (when it is considered in isolation). Similar inequalities can be obtained using monotonicity results [11, 8, 12] for other subnets. Throughout this paper, the transitions have single-server semantics. Our analysis can be extended immediately to STPN with bounded marking and in-nite-server transitions. Indeed, in such a case, each in-nite-server transition can be replaced by K single-server transitions, where K is the upper bound of the token numbers in the places. More precisely, we replace each in-nite-server transition t by K single-server transitions such a way that - p;t k An example of such a transformation is illustrated in Figure 7, where (a) (b)312 Figure 7: Transformation from in-nite-server transition to single-server transitions. (a) an in-nite-server transition. (b) equivalent single-server transitions. In our analysis, we assumed that the -ring times of each transition are i.i.d. exponential random variables with a -xed parameter. It is simple to extend the results to the case of marking dependent -ring rates, i.e., the -ring rate of a transition depends on the marking of input places, provided the number of dioeerent -ring rates is bounded. As an example, consider a transition t with a single input place. Let - 1 t and - 3 t be the -ring rates of INRIA Performance Analysis of Stochastic Timed Petri Nets 29 transition t when there are one token, two tokens, and more than 3 tokens in the input place. We replace the transition by three transitions t 1 , t 2 and t 3 with -ring rates - 1 t and - 3 respectively, in such a way that at any time at most one of the transitions is enabled, see Figure 8. The set of outgoing places are the same as that of t: oe t k 3. (a) (b)3123 Figure 8: Transformation from marking-dependent -ring rate to marking-independent -ring rates. (a) transition with 3 -ring rates. (b) equivalent transitions with -xed -ring rates. Our approach can be extended to the case that -ring times have phase-type distributions [37]. A phase-type distribution can be considered as the distribution of the time that a token passes through an ordinary Markovian state machine with a single source and a single sink transitions. The -ring times have exponential distributions for all transitions except the source and the sink which are immediate transitions. The sink transition represents the absorbing state. Let transition t have a phase-type distribution which is represented by a Markovian state machine N 0 with source transition t 0 and sink transition t s . We replace transition t in the original net by the subnet N 0 as follows. For any p 2 P and any transition t 0 6= t s of N 0 (including t 0 and excluding t s and oe t 0 . For any p 2 is a new place with initial marking 1. An example of the construction is illustrated in Figure 9 for an Erlang distribution with 3 stages. Recall that the -ring mechanism under consideration is race policy with age memory. The reader can therefore easily check that when transitions whose -ring times have phase-type distributions are thus replaced by corresponding Markovian state machines, we obtain a stochastically equivalent STPN with exponential -ring times. The performance measures considered in the paper are mostly the throughputs of transitions and the expectations of X p and X p e t . The same approach can be used to RR n\Sigma2642 Z. Liu (a)122224(b) Figure 9: Transformation from phase-type -ring times to exponential -ring times. (a) transition with 3-stages-Erlang-distribution -ring times. (b) transitions with exponential -ring times. INRIA Performance Analysis of Stochastic Timed Petri Nets 31 obtain linear equalities among higher moments of the token numbers. More precisely, for any m - 1, linear equalities can be established for variables E[X k ii . Similarly, for any m - 1, linear equalities can be established for variables E[X k 1 p2 (- n+1 )jF -n ii Finally, we remark that when the weights of the STPN are real numbers, all our analyses go through straightforwardly and the same results hold. Appendix Elimination of Immediate Transitions We present here a direct transformation technique which removes immediate transitions playing roles of synchronization and/or routing. We assume that for any such immediate transition t, t is the only output transition of all its input places, i.e. t. Further, we assume that for any immediate transition t, and ffl either oe t;p 0 ffl or oe t;p with a harmless abuse of notation, the index ffl p denotes the unique transition preceding place p. We show below that this kind of immediate transitions can be removed from the net without changing the -ring behavior of the other transitions. Consider a net E) with initial marking M and weights -; oe. Let t 0 be an immediate transition, g. Without loss of generality, we assume that min p2 ffl t 0 We construct a new net f E) with initial marking f M and weights ~ oe. The key idea is to create a place p i for each pair of input place p i and output place p j of transition t 0 . The set of input transitions of p i j is the union of ffl p i and ffl p j . The set of output transitions of p i . Such a transformation is illustrated in Figure 10, where transition t 0 is an immediate transition. The mathematical de-nition of f N is as follows. e RR n\Sigma2642 Z. Liu (a) (b) Figure 10: Removal of Immediate Transitions. (a): A subnet containing an immediate transition. (b): The subnet without immediate transition. e e where E is de-ned by The initial marking f M and weights ~ oe are de-ned accordingly: f ~ ~ oe t;p It is easily seen that if the sequences of the -rings times are the same for the same transitions in N and f -ring commencement and completion times are identical. The detailed proof can be done by induction and is left to the interested reader. INRIA Performance Analysis of Stochastic Timed Petri Nets 33 Acknowledgements : The author is very grateful to Dr Alain JEAN-MARIE for constructive comments and for eOEcient help in the computation of numerical results. --R iThe Eoeect of Execution Policies on the Semantics and Analysis of Stochastic Petri Netsj iA Class of Generalized Stochastic Petri Nets for the Performance Analysis of Multiprocessor Systemsj Performance Models of Multiprocessor Systems iErgodic Theory of Stochastic Petri Netsj iAnnotated Bibliography on Stochastic Petri Netsj Elements of Queueing Theory iParallel Simulation of Stochastic Petri Nets Using Recursive Equationsj iRecursive Equations and Basic properties of Timed Petri Netsj iStationary regime and stability of free-choice Petri netsj iEstimates of Cycle Times in Stochastic Petri Netsj iComparison Properties of Stochastic Decision Free Petri Netsj iGlobal and Local Monotonicities of Stochastic Petri Netsj iOptimization of Multiclass Queueing Networks: Polyhedral and Nonlinear Characterizations of Achievable Performancej Achievable Region and Side Constraintsj iA Characterisation of Independence for Competing Markov Chains with Applications to Stochastic Petri Netsj iA computational Algorithm for Closed Queueing Networks with Exponential Servers. iErgodicity and Throughput Bounds of Petri Nets with Unique Consistent Firing Count Vectorj iProperties and performance bounds for closed free choice synchronized monoclass queueing networksj iA General Iterative Technique for Approximate Throughput Computation of Stochastique Marked Graphsj iEmbedded Product-form Queueing Networks and the Improvement of Performance Bounds for Petri Net Systemsj iComputational Algorithms for Product Form Queueing Networks. iGeneralized Stochastic Petri Nets: A De-nition at the Net Level and Its Implicationsj iOperational Analysis of Timed Petri Nets and Application to the Computation of Performance Boundsj SPN: What is the Actual Role of Immediate Transitions? iSPNP: Stochastic Petri net packagej iExtended Stochastic Petri Nets: Applications and Analysisj iCyclic scheduling on parallel processors: an overviewj Markov Chain Models - Rarity and Exponentiality iPerformance Bounds for Queueing Networks and Scheduling Policiesj iStability of Queueing Networks and Scheduling Policiesj iComplete Decomposition of Stochastic Petri Nets Representing Generalized Service Networksj iPerformance Analysis using Stochastic Petri Netsj iQueueing Networks with Multiple Closed Chains: Theory and Computational Algorithms. iA Last Word on L --TR --CTR Jinjun Chen , Yun Yang, Adaptive selection of necessary and sufficient checkpoints for dynamic verification of temporal constraints in grid workflow systems, ACM Transactions on Autonomous and Adaptive Systems (TAAS), v.2 n.2, p.6-es, June 2007

mean token number;linear programming;throughput;stochastic timed petri net;uniformization;performance bound

291356

On Generalized Hamming Weights for Galois Ring Linear Codes.

The definition of generalized Hamming weights (GHW) for linear codes over Galois rings is discussed. The properties of GHW for Galois ring linear codes are stated. Upper and existence bounds for GHW of ZF_4 linear codes and a lower bound for GHW of the Kerdock code over Z_4 are derived. GHW of some ZF_4 linear codes are determined.

Introduction For any code D, -(D), the support of D, is the set of positions where not all the codewords of D are zero, and w s (D), the support weight of D, is the weight of -(D). For an [n; k] code C and any r, where 1 - r - k, the r-th Hamming weight is defined [7],[14] by d r is an [n; r] subcode of Cg: In [1],[4], and [5] different generalizations of GHW for nonlinear case were sug- gested. In the present paper we consider the case of Galois ring (GR) linear codes. Outline of the paper is following. In Section 2, we discuss the definition and state some properties of GHW of GR linear codes. In Section 3, we show that definition of GHW from [4] determines performance of group codes in the wire tap channel of type II. In Section 4, we find the number of different type subcodes in a Z 4 \Gammalinear code and state a connection between their support sizes. In Section 5, we use results of Section 4 to get bounds for GHW of Z 4 \Gammalinear codes. In Section 6, using results of sections 4 and 5, we obtain a lower bound for GHW of the Kerdock code over Z 4 and complete weight hierarchys of some short Z 4 \Gamma linear codes. We also show that though the minimum Hamming weight of a GR linear code can not exceed the minimum Hamming weight of an optimal linear codes, heigher weights of a GR linear code do can exceed corresponding weights of an optimal linear code of the same length and dimension. 2. Definition and Basic Properties Let R be a Galois ring, i.e., a finite commutative ring with identity e, whose set of zero divisors has the form pR for a certain prime p. After Nechaev's paper Galois rings have become easy to understand for a reader with a standard background in finite fields. In particular, using the definition given, one can prove [10] that and the characteristic of R (the order of e in the group (R; +)) equals p m . Since fixing the numbers p m and q m identifies R up to isomorphism, it may be also denoted as GR(q m of R form the following chain: and jR quotient ring is a Galois field of order q The ring R is constructed as a degree s Galois extension of Z p m in much the same way as one constructs the finite field from Z p . Note also that Galois rings encompass finite fields and residue rings as boundary cases. Namely, GF (p s Let C be a GR linear code of length n. Let D be any GR linear subcode of C. The support of D is defined exactly as for codes over a Galois field, i.e., the support of D is the set of not-all-zero symbol positions of D. In [1], the following definition was proposed for the r-th generalized Hamming weight of C d r linear subcode of C; log q One can see that this definition generalizes the definition of GHW for linear codes over Galois fields. L. A. Bassalygo proposed another definition for support and GHW [4]. Let A be any code of length n over arbitrary alphabet and let B be any subcode of A. Define the support of B as follows an such that a i 6= b i g: In other words, the support of B is the set of positions where not all the codewords of B have the same symbols. Define the following function is subcode of A and We will consider only those points where the function FA changes its values, that is, the following two sets are defined as follows and 1. It is clear that the function FA (M) is completely defined by the set of pairs so on. These pairs are called generalized Hamming weights. Example: Consider the nonlinear code This code has the following values and the function FA (M) can be depicted as follows (the function FA (M) is defined only for integer values of M , and we draw it everywhere just for visualization) 6 FA (M) One can see that Bassalygo's definition works for both linear and nonlinear cases. In the case of a linear code over GF (q), we have M i =M therefore we can study only values ffi i . In the general case, we must consider pairs In Corollary 1, we show that if C is a linear code over GR(q i.e., the function FC (M) can change its values only at see that in the GR linear case there is one-to-one correspondence between the set fd and the sets and knowledge of values d i allows us to find values vice versa. Note that for GHW of nonlinear codes one more definition, nonequivalent to Bassalygo's one, was given in [5]. Theorem 1 Let C be a linear code of length n over GR(q and if d r Proof: Let D be a GR linear subcode of C such that w s log q Dg. From linearity of D it follows that 0g. Then log q jD d r\Gammat (C) ! d r (C) for some t - m. To see that there exist a code A such that d r one can consider the code with generator matrix [1; Note that if C is a linear code of length n over GR(q code, then jCjjC mn. In the sequel, we suppose that ng is the set of positions of a code of length n. Denote by \Phi the direct sum of two sets. ng and C I \Phi fx Ig. I rg. Proof: Let S I I g, where x \Delta y denotes the inner product of x and y. Then log q jS I j log q jS ? I I be such that I I is a linear subcode of C, we see that d r (C) - jS I j - I rg. Let us establish inequality in the other direction. Let D be a linear subcode of C such that w s I rg. Theorem 2 Let C be a GR linear code and let C ? be its dual code. Then fd r times times times Proof: It suffices to prove that if then there are no more than m\Gammal values d t+1 that d t (C) ! d t+1 (C); d t+m\Gammal (C) ! d t+m\Gammal+1 (C), and d t+1 for some t. At first we show that d t (C) - ng n -(D) I \Phi D. Hence, jC ? I 1, we obtain d t (C) - jI log Next we show that if d a subcode of C such that log q ng n -(D) and Hence jC I By Lemma 1, we obtain d where contradiction. Corollary 2 Let g. Then ng l log q M i =M there is not ffi j such that ffi ? Good nonlinear codes can be constructed by mapping GR linear codes into codes over GF (q) [10],[6]. Therefore it is natural to ask whether GR linear codes themselves can have better characteristics than linear codes over Galois fields. The following theorem shows that GR linear codes can not be better (from the point of view of length, cardinality, and minimum Hamming distance) than optimal linear codes, i.e. than linear codes having the largest cardinality for a given length and distance. Theorem 3 Let C be a GR linear code over GR(q Hamming distance d. Then there exist a linear code D over GF (q m ) of length n with minimum distance d such that jDj - jCj. Proof: We consider the case of a code C over GR(q The case of an arbitrary m can be proved by the same way. A generator matrix of C can be written in the are r 1 \Theta n and r 2 \Theta n GR(q C be a subcode of C with the generator matrix Obviously, the code b C has the same length and the same minimum distance as C. Note that b C contains only elements from the ideal pGR(q us define exterior multiplication of an element ff 2 pGR(q by an element a 2 GR(q Analogously, aff n ). Now taking into account (1), it is easy to see that b C is a vector space over GR(q of dimension r Hence b C is isomorphic to some [n; r over GF (q). Let D be a code with a generator matrix of the code b considered as a matrix over GF (q 2 ). Then D is a linear code over GF (q 2 ) of length n with minimum distance d and Using the Nordstrom-Robinson code over Z 4 as an example, we shall show later that a GR linear code can have the same length, cardinality, and minimum distance as an optimal linear code, and the function F for such a GR linear code can exceed the function F for the optimal linear code. 3. An Application to The Wire-Tap Channel of Type II One of motivations for the introduction of a notion of GHW was cryptographical one. Ozarow and Wyner [11] considered using linear codes on the wire-tap channel of Type II. One of their schemes uses an [n; k] linear code, say C, over GF (q). The code has q n\Gammak cosets, each representing a q-ary (n \Gamma k)-tuple. If the sender wants to q\Gammaary symbols of information to the receiver, he selects a random vector in the corresponding coset. The adversary has full knowledge of the code, but not of the random selection of a vector in a coset. Wei showed [14] that in the linear case if the adversary is allowed to tap s symbols (of his choice) from the sender, he will obtain r q\Gammaary symbols of information, if and only if s - d r (C ? It is obvious that we can use a group code on the said channel. Let C be a group code of length n over an additive group \Theta; '. The code C has ' n =jCj cosets. As in the linear case, the number of a coset is transmitted information and the sender transmits a random vector from the corresponding coset. Let C have values let the code C be used on the wire-tap channel of type II. We prove the following theorem. Theorem 4 Let s symbols be taped from transmitter and the number of cosets of C that can (equiprobably) contain the transmitted vector (with known s taped symbols) is equal to or greater than ' n\Gammas Proof: Let I be a set of taped positions. By L denote a set of vectors that contain zeros on positions from I , i.e., Ig. Cosets of C that contain vectors from L form a subgroup, say \Psi, of the quotient group of cosets of C. Cosets of C that contain vectors with given values on positions from I form a coset of \Psi. Hence w.l.o.g. we can assume that taped symbols are zeros. Denote by D a set of codewords of C that belong to L. It is clear that any coset of C that belongs to \Psi has exactly jDj vectors from L and hence The best strategy for the adversary is to choose the set I such that the cardinality of D would be maximal. We claim that jDj - M i . Indeed, suppose that there exists a subcode B ' C such that jBj ? M i and the definition of GHW the support weight of any subcode of cardinality larger than M i is equal to or greater than ffi i+1 . A contradiction. Hence the transmitted vector equiprobably belongs to one of j\Psij cosets of C, and j\Psij - ' n\Gammas 4. Subcodes of a Z 4 \Gammalinear code Let C be a Z 4 \Gammalinear code of length n and cardinality 4 r1 2 r2 with a generator matrix are r 1 \Theta n and r 2 \Theta n Z 4 \Gammamatrices. We will say that C is an C denote a subcode of C with the generator matrix [2G 1 ] and by e C denote a subcode of C with the generator matrix Remark. Throughout the rest of the paper we use symbols - and - in this meaning Note that b C and e C are [n; 0; r 1 ] and [n; 0; r an [n; s of C and let be a generator matrix of D, where G 1 and G 2 are s 1 \Theta n and s 2 \Theta n Z 4 \Gammamatrices respectively. By denote the set of subcodes [n; s of C. In the sequel, we need the following propositions. Proposition 1 Proof: Consider the number of ways of choosing s 1 Z 4 \Gammalinearly independent codewords. Note that linear independence over Z 4 implies that none of these codewords consists only of zero divisors of Z 4 . The total number of codewords of C of them consist of zero divisors of Z 4 . Hence the first codeword, say u 1 , can be chosen by 2 distinct ways. The second codeword u 2 can not be equal to au 1 C . Indeed, if then either 2u are not Z 4 \Gammalinearly independent. Thus u 2 can be chosen by 2 ways. Analogously, if we have codewords then a codeword u t can be chosen by 2 distinct ways. Thus the number of ways of choosing s 1 Z 4 \Gammalinearly independent codewords equals Now we have to choose s 2 Z 2 \Gammalinearly independent codewords that consist from zero divisors of Z 4 . These codewords should also be Z 2 \Gammalinearly independent from the s 1 codewords chosen earlier. The number of ways for choosing these codewords equals Using the same arguments, one can see that the number of distinct generator matrices of any [n; s To complete the proof, we multiply (2) by (3) and divide by (4). In the linear case, any nonzero codeword of a linear code belongs to one and the same number of subcodes of a given dimension. The situation is different in the Z 4 \Gammalinear case. We should consider the following three cases. Proposition 2 Let u 2 C; 2u 6= 0. Then u belongs to subcodes from -(C; s Proof: Recall that the generator matrix of C has the form . To prove the proposition, we assume that one of rows of the matrix G 1 equals u and use arguments from the proof of Proposition 1. Proposition 3 Let subcodes from -(C; s Proof: Suppose a codeword u 2 D; then either D or u 2 e D. Consider the case u 2 b D. The subcode D contains 2 w such that w 2 b D. The code C contains 2 codewords v such that C. Hence the number of subcodes D such that u 2 b D equals the total number of [n; s subcodes of C multiplied by the factor 2 Using Proposition 1, we get Consider the case u 2 e D. From (5) it follows that the number of subcodes D such that u 62 b Assuming that one of rows of G 2 equals u and using arguments like ones from Proposition 1, we get the number of subcodes D such that u 2 e D. This number equals Summation of (5) and (6) completes the proof. Proposition 4 Let u 2 e C. Then u belongs to subcodes from -(C; s Proof: Suppose that u 2 D. From the conditions of the proposition it follows that u 62 b D. Hence the matrix G 1 can be formed by any s 1 Z 4 \Gammalinearly independent codewords of C. Assuming that one of rows of the matrix G 2 equals u and using arguments from the proof of Proposition 1, we complete the proof. By wtL (u) denote the Lee weight of a vector u. Proposition 5 ([3], [13]) For an [n; r Using an approach suggested in [8], we can obtain a connection between support weight of C and supports weights of its subcodes. Proposition 6 C) C) are numbers from Propositions 2,3, and 4. Proof: Using Proposition 5 and Propositions 2,3, and 4, we get (an abbreviation s.t. means such u2b Cn0 u2e Cnb C u2b Cn0 u2e Cnb C u2b Cn0 u2e Cnb C wtL (u)C A u2b Cn0 u2e Cnb C wtL (u)C A u2b Cn0 C) It is known that if C is a Z 4 \Gammalinear code and C is its binary image under the Gray map, then minimum Lee distance of C equals minimum Hamming distance of C [6]. Such good nonlinear binary codes as Kerdock, Delsarte-Goethals, and Preparata codes can be constructed by Gray mapping Z 4 \Gammalinear codes [6]. Therefore it is natural to ask what is the minimal length, N(r of an [n; r code with minimum Lee distance dL . We confine ourselves to the case of [n; codes and obtain an analog of the Griesmer bound for N(r; 0; dL ). Theorem 5 dL Proof: Let C be an [n; \Gammalinear code with minimum Lee distance dL . Note that in this case b Using Proposition 6, we get Obviously, w s ( b Assuming w s ( b From (7) and (8) it follows that If 3ag is an [n; \Upsilon . Otherwise wtL (2a) ! dL . Substituting the value for N(1; 0; dL ) in (9) completes the proof. Note that the Nordstrom-Robinson code over Z 4 meets this bound. 5. Bounds on Higher Weights We start with Plotkin type bound. Let d s1 ;s 2 is an [n; s subcode of Cg. Theorem 6 If C is an [n; r d 2s (C) - n Proof: Calculate in two ways the sum, say S, of codeword Lee weights of all subcodes of C. The total number of subcodes equals -(r According to Proposition 5 the sum of Lee weights of codewords of any [n; s subcode is greater than or equal to 2 2s1+s2 d 2s1 +s2 (C). Hence Let A be a 2 2r1+r2 \Theta n matrix whose rows are codewords of C. W.l.o.g. we can consider the first column of A. Any a 2 Z 4 occurs 2 2r1+r2 \Gamma2 times in the column. belongs to A 1 [n; s of C and there are 2 \Delta 2 2r1+r2 \Gamma2 such vectors. If a = 2 then there are three different cases. C. Then v belongs to A 1 [n; s subcodes of C and there are C. Then v belongs to A 2 [n; s subcodes of C and there are 2 r1 \Gamma1 such vectors. C. Then v belongs to subcofdes of C and there are 2 r1 +r2 Hence After computations we get Setting respectively, we get the assertion. Let D be a Z 4 \Gammalinear code and C be its [n; r subcode such that w s d r1 ;r 2 (D). Let fl r1 ;r 2 and fl r1 be upper bounds for support sizes of subcodes e C and C. Remark. If we don't have an information on the structure of subcodes e D and D then we can say that (D). Sometimes, if we know weight distribution of support sizes of all subcodes of codes e D and b D, we can find exact values for support sizes of subcodes e C and b C. (see next chapter). Now we can get a Griesmer type bound. Theorem 7 If 2 r1 +r2 \Gammas 1 d r1 ;r 2 (D) - otherwise d r1 ;r 2 d r1 ;r 2 (D) Proof: From proposition 6 we have (D) C) C) C) C) It is easy to check that A 2 \GammaA 3 is always nonnegative when s Replacing w s ( b C) by fl r1 and w s ( e C) by fl r1 ;r 2 (C) in the case when A by d r1 ;r 2 (D) otherwise, we get the assertion. As corollarys we have the following relations. Although in general d r can be equal to d r+1 , d s1 ;s 2 +1 is always greater than d s1 ;s 2 Corollary 3 Proof: Inequality d s1 ;s 2 +1 (D) - d 0;s1 +s2+1 (D) is obvious. Assuming r (like in Theorem 7) D) For linear binary codes we have [8] s D and b are isomorphic to linear binary codes we have D) D) - d 0;s1 (D) the assertion follows. Corollary 4 ae oe Proof: If r 1) and the assertion follows from assuming fl r1 ;r 2 (D) and Theorem 7. Note that as an estimation for value d 0;t (D) we can use any lower bound for generalized Hamming weights of binary liner codes. For example using the Griesmer bound we get d s1 ;0 (D); where d(C) is minimum Hamming weight of C. Note that some codes have equalities in (10) and (11) (see Remark 5). Our last bound is an existence one. Theorem 8 For any s 1 - r 1 and s there exists an [n; r Z 4 \Gammalinear code with d s1 ;s 2 - d which satisfies d Proof: The number of [n; r codes that contain a given subcode [n; s equals d is the number of [n; s codes with supports of size less than or equal to d. Hence if d is less than the total number of [n; r at least one of [n; r codes has a subcode [n; s greater than or equal to d. Substitution valies for the function - gives us the assertion. 6. Determination of Weight Hierarchy It is often that we know more about subcodes e C and b C, wich are isomorphic to linear binary codes. In those cases we can get better bounds. Consider the Kerdock code over Z 4 [6],[10]. In our notation the Kerdock code Km is an [2 code, m is odd. The minimum Lee distance of Km is and the minimum Hamming distance is 2 . Moreover, in [6] compositions of all codewords and weight distribution of Km were found. Suppose l t Km such that -(u) 6= has one of the following compositions l l be the number of codewords of weight i. Then In [6], it was also shown that subcode b Km is isomorphic to the first order Reed-Muller code RM(1;m) of length 2 m . To find a lower bound for d r (Km ) we need the following lemma. Lemma 2 The support size of any r dimensional subcode, say C, of RM(1;m) equals to either Proof: If C contains the all-one codeword then w s that C does not contain the all-one codeword. It is known [14](Theorem 5) that d r any r rows (that are not all-one) in a standard generator matrix generate a subcode whose support size equals d r (RM(1;m)). Any codeword of RM(1;m) can be consedered as a boolean linear function of m variables, say v affine transformation of v belongs to the automorphism group of RM(1;m) and hence it does not change the support size of C. Let f 1 functions corresponding to basis vectors of C. Assume that f 1 (v) has a term v 1 . Then the affine transform v 1 to the form v 1 and f 2 (v) to some f 0 (v). Since f 2 (v) is linearly independent of f 1 (v) and f 1 (v) 2 (v) has a term v j The affine transform v 2 2 (v) to the form v 2 . Continuing this procedure, we get f 1 that is r rows from a standard generator matrix. Theorem 9 d d r d we have (like in Theorem 7) where C is an [n; r of Km such that w s C and e C are [n; 0; r 1 ] and [n; 0; r subcodes of Km respectively, and they are isomorphic to [n; r 1 ] and [n; r subcodes of RM(1;m). Hence from the Lemma 2 we get From possible compositions of codewords of Km it follows that if V is an [n; 1; 0] subcode of Km then w s (V Hence (3 By denote the right side of (14). It is evident that d r (Km It is easy to check that n(r 1). On the other hand, since w s (C) - w s ( e C), we have linear subcode of Km ; ae d r=2 (RM(1;m); if r is even, d (r+1)=2 (RM(1;m); if r is odd: It is easy to check that d r (RM(1;m)) at r - (m \Gamma 1)=2. It is obvious that for Km Indeed, it is enough to take a subcode C =! a ?= f0; a; 2a; 3ag generated by a codeword a with the composition l 0 (note that in Km there are also subcodes of cardinality 4 generated by a couple of codewords consisting of zero divisors of Z 4 ), and check that support weight of C meets the lower bound and The Kerdock code over Z 4 can be also defined for the case of even m [6]. In this case Similar to the case when m is odd we get Theorem d d Remark. Theorems 9 and 10 were obtained for the first time in [2] [3]. Later and independently the following estimates were obtained for the Kerdock codes [13]: d ae P 2r and d It is easy to check that estimates (16)and (17) coincides with estimates from Theorems m=2. For the r - m=2 estimates from the theorems are better than estimates (16) and (17). Minimum distance of a code is only a small part of information contained in the spectrum of distance distribution of the code. In the same way, the number of subcodes of a given dimension with a given support weight gives much more information on a code than only minimum support weights of subcodes. For example we used distribution of support weights of the first order Reed-Muller code in the proof of Theorem 9 and 10. Papers [7],[12] are devoted to studying distributions of support weights of linear codes. We shall find weight distribution of Z 4 \Gammalinear sub-codes of cardinality 2 3 (weight distribution of subcodes of cardinality 2 2 is evident). Let A 3 linear subcode of Km ; Theorem 11 A 3 A 3 A 3 Proof: We should consider subcodes Consider an [2 generated by codewords a; 2a 6= 0; -(a) 6= Assuming using Proposition 6, we get that and w s 3 be sets of [2 subcodes with support weights 7 It is easy to see that if w s (! a; b Hence, As it was mentioned if w s any codeword a 2 C; 2a 6= 0, has Hamming weight 3 . Hence the total number of codewords of weight 3 in all the subcodes from S 1 equals In Km there are 2 m+1 \Gamma4 codewords b; l 0 , that are Z 4 \Gammalinearly independent form a given a 2 Km ; 2a 6= 0. According to (18), half of these codewords half of them form subcodes So, a given codeword a; 2a 6= 0, belongs to 2 subcodes from S 1 . Hence the the total number of codewords of weight 3 in all the subcodes from S 1 equals the half number of codewords of this weight (codewords a and 3a always belong to one and the same subcode) multiplied by 2. So, using (19) and (12), we get subcode of Km with support contains the codeword Assuming s using Proposition 3, we get that the number of an [2 subcodes with support According to Proposition 1 the total number of an [2 subcodes equals 1). Hence the number of [2 subcodes with support weight 7 Since 2Km is isomorphic to RM(1;m), one can see that in 2Km there are \Theta m subcodes of Km with support weight 2 m and \Theta m+1\Gamma with support weight 7 The numbers jS 1 \Theta m+1\Gamma \Theta m; jS 3 j, and A 2 \Theta mgive the numbers of Z 4 \Gammalinear subcodes of cardinality 8 with support weights 7 Corollary 5 The values ffi 3 and M 3 for the Kerdock code Km are Using Theorems 2 and 7, we can find the weight hierarchy of some short Z 4 -linear codes. Theorem 12 The Nordstrom-Robinson code (NR) over Z 4 has weight hierarchy Proof: The Nordstrom-Robinson code is a self-dual code with d 1 According to (15), d 2 5. From this fact and Theorem 2 it follows that 8g. The claim of the theorem follows from this hierarchy. Remark. The optimal linear code, say B, over GF (4) of length 8 and dimension 4 has 4: The best possible function FB (M) for such a code and the function FNR (M) have the following form (see the picture). One can see that 5. Thus the Nordstrom-Robinson over Z 4 code is better, in some sense, than the optimal linear code over GF (4) of the same length and dimension. From cryptographical point of view it means that if we use the code B on the wire tap channel the adversary must tap 4 symbols from transmitter to get 2 bits of information, whereas in the case of using the Nordstrom-Robinson code the adversary must tap at least 5 symbols to get 2 bits of information. Whether there exist other GR linear codes (or codes over another ring or additive group) that are better (from the point of view of GHW) than optimal linear codes over GF (q) is an open question. Theorem 13 The Kerdock code of length 16 over Z 4 has weight hierarchy Proof: From (15) and Corollary 5 it follows that fd 1 (K 4 5 be the Preparata code over Z 4 of length 16. Using the 6 F(M) 4\GammaThe Nordstrom-Robinson code over Z 4 3\GammaThe optimal [8; 4] linear code over GF (4) MacWilliams identities for Z 4 \Gammalinear codes [6], we see that in P 5 there exists a codeword a such that 2a 6= 0 and 4. Using this fact and taking into account that minimum Hamming weight of P 4 equals 4, we get d 1 Now using Theorem 2, we get the hierarchy of K 4 . Remark. Note that if we start with known values d 1 (Km ) and d 2 (Km then in (10) and (11) we have equality for all d s1 ;0 and d s1 ;1 . Using Theorems 2 and 13, we can also find weight hierarchy of the Preparata code P 4 Theorem 14 The Kerdock code of length 32 over Z 4 has weight hierarchy Proof: From (15) and Theorem 7 it follows that fd 1 (K 5 5 be the Preparata code over Z 4 of length 32. Using the MacWilliams identities for Z 4 \Gammalinear codes [6], we see that in P 5 there exists a codeword a such that 2a 6= 0 and 5. Using this fact and taking into account that minimum Hamming weight of P 5 equals 4, we get d 1 Now using Theorem 2, we get fd 5 (K 5 From Theorem 6 it follows that d 4 (K 5 or 28. Suppose d 4 (K 5 contradiction. Thus d 4 (K 5 27. The claim of the theorem follows from the hierarchy of the values d i (K 5 ). Using Theorems 9 and 2, we can also find weight hierarchy of the Preparata code --R "Generalized Hamming Weights for Z4 -Linear Codes," "On Generalized Hamming Weights for Galois Ring Linear Codes," "On Generalized Hamming Weights for Galois Ring Linear Codes," "Supports of a Code," "Upper Bounds on Generalized Distances," "The Z4 - Linearity of Kerdock, Preparata, Goethals and Related Codes," "The Weight Distribution of Irreducible Cyclic Codes with Block Lengths n 1 ((q l \Gamma 1)=N )," "Generalization of the Griesmer Bound," The Theory of Error-Correcting Codes "Kerdock Code in a Cyclic Form," "Wire-Tap Channel II," "The Effective Length of Subcodes," "On the Weight Hierarchy of the Kerdock Codes Over Z 4 ," "Generalized Hamming Weights for Linear Codes," --TR --CTR Cui Jie, Support weight distribution of Z Manish K. Gupta , Mahesh C. Bhandari , Arbind K. Lal, On Linear Codes over$${\mathbb{Z}}_{2^{s}}$$, Designs, Codes and Cryptography, v.36 n.3, p.227-244, September 2005 Keisuke Shiromoto , Leo Storme, A Griesmer bound for linear codes over finite quasi-Frobenius rings, Discrete Applied Mathematics, v.128 n.1, p.263-274, 15 May

generalized Hamming weights;galois ring linear codes

291364

Time- and VLSI-Optimal Sorting on Enhanced Meshes.

AbstractSorting is a fundamental problem with applications in all areas of computer science and engineering. In this work, we address the problem of sorting on mesh connected computers enhanced by endowing each row and each column with its own dedicated high-speed bus. This architecture, commonly referred to as a mesh with multiple broadcasting, is commercially available and has been adopted by the DAP family of multiprocessors. Somewhat surprisingly, the problem of sorting m, (mn), elements on a mesh with multiple broadcasting of size $\sqrt n\times \sqrt n$ has been studied, thus far, only in the sparse case, where $m\in \Theta \left( {\sqrt n} \right)$ and in the dense case, where m(n). Yet, many applications require using an existing platform of size $\sqrt n\times \sqrt n$ for sorting m elements, with $\sqrt n

Introduction With the tremendous advances in VLSI, it is technologically feasible and economically viable to build parallel machines featuring tens of thousands of processors [4, 22, 39, 40, 42, 47]. However, practice indicates that this increase in raw computational power does not always translate into performance increases of the same order of magnitude. The reason seems to be twofold: first, not all problems are known to admit efficient parallel solutions; second, parallel computation requires interprocessor communication which often acts as a bottleneck in parallel machines. In this context, the mesh has emerged as one of the platforms of choice for solving problems in image processing, computer vision, pattern recognition, robotics, and computational morphology, with the number of application domains that benefit from this simple and intuitive architecture growing by the day [2, 4, 39, 40, 42]. Its regular and intuitive topology makes the mesh eminently suitable for VLSI implementation, with several models built over the years. Examples include the ILLIAC V, the STARAN, the MPP, and the MasPar, among many others [4, 5, 11, 39]. Yet, the mesh is not for everyone: its large computational diameter makes the mesh notoriously slow in contexts where the computation involves global data movement. To address this shortcoming, the mesh has been enhanced by the addition of various types of bus systems [1, 2, 16, 18, 20, 23, 39, 43]. Early solutions involving the addition of one or more global buses shared by all the processors in the mesh, have been implemented on a number of massively parallel machines [2, 4, 19, 39]. Yet another popular way of enhancing the mesh architecture involves endowing every row with its own bus. The resulting architecture is referred to as mesh with row buses and has received a good deal of attention in the literature. Recently, a more powerful architecture has been obtained by adding one bus to every row and to every column in the mesh, as illustrated in Figure 1. In [20] an abstraction of such a system is referred to as mesh with multiple broadcasting, (MMB, for short). The MMB has been implemented in VLSI and is commercially available in the DAP family of multicomputers [20, 38, 41]. In turn, due to its commercial availability, the MMB has attracted a great deal of attention. Applications ranging from image processing [21, 38, 41], to visibility and robotics [7, 14, 15, 37], to pattern recognition [9, 10, 20, 25, 37], to optimization [17], to query processing [9, 13], and to other fundamental problems [3, 6, 8, 9, 10, 16] have found efficient solutions on this platform and some of its variants [18, 25, 29, 30]. As we shall discuss in Section 2, we assume that the MMB communicates with the outside world via I/O ports placed along the leftmost column of the platform. This is consistent with the view that enhanced meshes can serve as fast dedicated coprocessors for present-day computers [28]. In such a scenario, the host computer passes data, in a systolic fashion, to the enhanced mesh in batches of n and so, in the presence of m elements to be processed, the leftmost m columns will be used to store the input. From now on, we will not be concerned with I/O issues, assuming that the input has been pretiled in the leftmost columns of the platform. Sorting is unquestionably one of the most extensively investigated topics in computer science. Somewhat surprisingly, thus far, the problem of sorting m, (m - n), elements on an MMB of size addressed only in the sparse case, where m 2 \Theta( n) or in the dense case, when m 2 \Theta(n). For the sparse case, Olariu et al. [35] have proposed a time-optimal \Theta(log n) time algorithm to sort elements stored in one row of an MMB of size n. Yet, many applications require using an existing platform of size for sorting m elements, with For example, in automated visual inspection one is interested in computing a similarity measure coefficient for a vertical strip of the image [27]. A similar situation arises in the task of tracking a mobile target across a series of frames [31, 46]. In this latter case, the scene consists typically of static objects and one is interested only in evaluating movement-related parameters of the target. In order to perform this task in real-time it is crucial to use the existing platform to process (sort, in parameter values at pixels in a narrow rectangular subimage only. The width of the subimage of interest depends, of course, on the speed with which the target moves across the domain. Therefore, in order to be able to meet read-time requirements one has to use adaptive sorting algorithms that are as fast as possible. The main contribution of this paper is to propose the first know adaptive time- and VLSI-optimal sorting algorithm on the MMB. Specifically, we show that once we fix a pretiled in the leftmost m columns of an MMB of size time. We show that this is time-optimal for this architecture. It is also easy to see that this achieves the AT 2 lower bound in the word model. At the heart of our algorithms lies a novel deterministic sampling scheme reminiscent of the one developed recently by Olariu and Schwing [36]. The main feature of our sampling scheme is that, when used for bucket sorting, the resulting buckets are well balanced, making costly rebalancing unnecessary. The remainder of this paper is organized as follows: Section 2 discusses the model of computation used throughout the paper; Section 3 presents our time- and VLSI lower bound arguments; Section 4 reviews a number of basic data movement results; Section 5 presents the details of our optimal sorting algorithm. Finally, Section 6 summarizes the results and poses a number of open problems. 2 The Model of Computation An MMB of size M \Theta N , hereafter referred to as a mesh when no confusion is possible, consists of MN identical processors positioned on a rectangular array overlaid with a high-speed bus system. In every row of the mesh the processors are connected to a horizontal bus; similarly, in every column the processors are connected to a vertical bus as illustrated in Figure 1. We assume that the processors in the leftmost column serve as I/O ports, as illustrated, and that this is the only way the MMB can communicate with the outside world. Figure 1: An MMB of size 4 \Theta 4 Processor P (i; j) is located in row i and column j, (1 - in the northwest corner of the mesh. Each processor P (i; j) has local links to its neighbors exist. Throughout this paper, we assume that the MMB operates in SIMD mode: in each time unit, the same instruction is broadcast to all processors, which execute it and wait for the next instruction. Each processor is assumed to know its own coordinates within the mesh and to have a constant number of registers of size O(log MN ); in unit time, the processors perform some arithmetic or boolean operation, communicate with one of their neighbors using a local link, broadcast a value on a bus, or read a value from a specified bus. Each of these operations involves handling at most O(log MN) bits of information. Due to physical constraints, only one processor is allowed to broadcast on a given bus at any one time. By contrast, all the processors on the bus can simultaneously read the value being broadcast. In accord with other researchers [3, 16, 17, 20, 21, 23, 38, 39], we assume constant broadcast delay. Although inexact, experiments with the DAP, the PPA, and the YUPPIE multiprocessor array systems seem to indicate that this is a reasonable working hypothesis [23, 26, 29, 38, 39]. 3 The Lower Bounds The purpose of this section is to show that every algorithm that sorts m; m - n, elements pretiled in the leftmost d m e columns of an MMB of size must take time. Our argument is of information transfer type. Consider the submesh M consisting of processors n . The input will be constructed in such a way that every element initially input into M will find its final position in the sorted order outside of M. To see that this is possible, note that m - n guarantees that the number of elements in M satisfies: mp - mSince at most O( m ) elements can leave or enter M in O(1) time, it follows that any algorithm that correctly sorts the input data must take at time. Thus, we have the following result. Theorem 3.1. Every algorithm that sorts m, (m - n), elements stored in the leftmost m columns of a mesh with multiple broadcasting of size must take In addition to time-lower bounds for algorithms solving a given problem, one is often interested in designing algorithms that feature a good VLSI performance. One of the most m//n m/2/n Figure 2: Illustrating the lower bound argument commonly used metrics for assessing the goodness of an algorithm implemented in VLSI is the AT 2 complexity [44], where A is the chip area and T is the computation time. A time lower bound based on this metric is strong because it is not based on memory requirements or input/output rate, but on the requirements for information flow within the chip. It is well-known [44] that in the word model the AT 2 lower bound for sorting m elements on a VLSI chip is m 2 . In our case, the size m of the input varies, while the area, n, of the mesh is a constant. Hence, for any algorithm to be AT 2 -optimal, we must have T is the running time. Thus, in this case, the time lower bound of \Omega\Gamma m into VLSI-optimality. In the remaining part of the paper, we show that the lower bounds derived above are tight by providing an algorithm with a matching running time. For sake of better understanding, we first present a preliminary discussion on some data movement techniques used throughout the paper. 4 Data Movement Data movement operations constitute the basic building blocks that lay the foundations of many efficient algorithms for parallel machines constructed as an interconnection network of processors. The purpose of this section is to review a number of data movement techniques for the MMB that will be instrumental in the design of our sorting algorithm. Merging two sorted sequences is one of the fundamental operations in computer science. Olariu et al. [35] have proposed a constant time algorithm to merge two sorted sequences of total length stored in one row of an MMB of size precisely, the following result was established in [35]. Proposition 4.1. Let S p n, be sorted sequences stored in the first row of an MMB of size holding a s). The two sequences can be merged into a sorted sequence in O(1) time. Since merging is an important ingredient in our algorithm, we now give the details of the merging algorithm [35]. To begin, using vertical buses, the first row is replicated in all rows of the mesh. Next, in every row i, (1 - i - r), processor P (i; i) broadcasts a i horizontally on the corresponding row bus. It is easy to see that for every i, a unique processor P (i; j), . Clearly, this unique processor can now use the horizontal bus to broadcast j back to P (i; i). In turn, this processor has enough information to compute the position of a i in the sorted sequence. In exactly the same way, the position of every b j in the sorted sequence can be computed in O(1) time. Knowing their positions in the sorted sequence, the elements can be moved to their final positions in time. Next, we consider the problem of merging multiple sorted sequences with a common length. Let a sequence of n be stored, one per processor, in the first row of an MMB of size n. Suppose that the sequence consists of k sorted subsequences and each subsequence consists of consecutive elements of the original sequence. The goal is to sort the entire sequence. For definiteness, we assume that subsequence j, (1 - j - k), contains the elements a (j \Gamma1) . Our sorting algorithm proceeds by merging the subsequences two at a time into longer and longer subsequences. The details are as follows. We set aside submeshes of size 2 \Theta 2 k and every pair of consecutive subsequences is merged in each one of these submeshes. Specifically, the first pair of subsequences is allocated the submesh with its north-west corner; the next pair of subsequences is allocated the submesh with processor its north-west corner, and so on. Note that moving the subsequences to the corresponding submeshes amounts to a simple broadcast operation on vertical buses. Now in each submesh, the corresponding subsequences are merged using the algorithm described above. By Proposition 4.1, this operation takes constant time. By repeating the merging operation dlog ke times, the entire sequence is sorted. Consequently, we have the following result. Lemma 4.2. A sequence consisting of k equal-sized sorted subsequences stored in the first row of a mesh with multiple broadcasting of size can be sorted in O(log Finally, we look at a data movement technique on an MMB of size involving the reorganization of the elements in the leftmost x columns of the mesh sorted in row-major order to column-major order (see Figure 3(a) to 3(d)). This can be accomplished by a series of simple data movement operations whose details follow. To simplify the notation, we shall assume that x is an integer. The leftmost x columns of the mesh are moved, one at a time, as follows. For each column s being moved, every processor P broadcasts the element it holds to processor P (r; illustrated in Figure 3(b). We now view the mesh as consisting of horizontal submeshes R 1 each of size x \Theta p n. In a submesh R p , x broadcasts its value along column bus j and P (j; records it as shown in Figure 3(c). Again, in constant time, each processor P (j; broadcasts its value along row bus j to processor P (p; j). The above can be repeated for each submesh R p with 1 - p - x, thus accomplishing the required data movement in O(x) time. To summarize our findings we state the following result. Lemma 4.3. Given a mesh with multiple broadcasting of size stored in the leftmost x columns in sorted row-major order, the data can be moved into sorted column-major order in the leftmost x columns in O(x) time. (a) (b) (c) (d) R x x x Figure 3: Illustrating the data movement of Lemma 4.3 5 The Algorithm We are now in a position to present our time- and VLSI-optimal sorting algorithm for the MMB. Essentially, our algorithm implements the well-known bucket sort strategy. The novelty of our approach resides in the way we define the buckets, ensuring that no bucket is overly full. Throughout, we assume an MMB R of size Fix an arbitrary constant 1. The input is assumed to be a set S of m elements from a totally ordered universe 1 stored in the leftmost d m e columns of R. To We assume O(1) time comparisons among the elements in the universe. avoid tedious, but otherwise inconsequential, details we assume that m is an integer. The goal is to sort these elements in column-major order, so that they can be output from the mesh in O( m time. We propose to show that with the above assumptions the entire task of sorting can be performed in O( m time. Thus, from our discussion in Section 3, we can conclude that our algorithm is both time- and VLSI-optimal. It is worth mentioning yet another interesting feature of our algorithm, namely, that the time to input the data, the time to sort, and the time to output the data are essentially the same. To make the presentation more transparent and easier to follow we refer to the submesh consisting of the leftmost m columns of R as M. In other words, M is the submesh that initially contains the input. Further, a slice of size k of the input consists of the elements stored in k consecutive rows of M. We will first present an outline of our algorithm and then proceed with the details. We start by partitioning M into slices of size m n and sort the elements in each such slice in row-major in O( m using an optimal sorting algorithm for meshes [32, 45]. Next, we use bucketsort to merge consecutive m of these into slices of size ( m order. Using the same strategy, these slices are again merged into larger slices sorted in row-major order. We continue the merging process until we are left with one slice of size row-major order. Finally, employing the data movement discussed in Lemma 4.3, the data is converted into column-major order. We proceed to show that the task of merging m consecutive sorted slices of size ( m into a sorted slice of size ( m time. For this purpose, it is convenient to view the original mesh R as consisting of submeshes R j;k of size ( m involving processors P (r; s) such that (j \Gamma 1)( m We refer to submeshes R k;k as diagonal - see Figure 4 for an illustration. Notice that the diagonal submeshes can be viewed as independent MMBs, since the same task can be performed, in parallel, in all of them without broadcasting conflict. The algorithm begins by moving the elements in every R k;1 to the diagonal submesh R k;k . This can be accomplished, column by column, in O( m time. We now present the details of the processing that takes place in parallel in every diagonal submesh R k;k . The rightmost element in every row of R k;k will be referred to as the leader of that row as shown in Figure 4. To begin, the sequence of leaders in increasing order. Note that by virtue of our grouping, the sequence of leaders consists leaders R R 11 R 22 R 22 Figure 4: Illustrating diagonal submeshes and leaders of m sorted subsequences, and so, by Lemma 4.2, the sequence of leaders can be sorted in O(log m time. Let this sorted sequence be a 1 , a For convenience, we assign a Next, in preparation for bucket sort, we define a set of ( m such that for every j, (1 (2) By definition, the leaders a (j \Gamma1)m p n +1 through a jm belong to bucket B j . This observation motivates us to call a row in R k;k regular with respect to bucket B j if its leader belongs to B j . Similarly, a row of R k;k is said to be special with respect to bucket B j if its leader belongs to a bucket B t with t ? j, while the leader of the previous row belongs to a bucket To handle the boundary case, we also say that a row is special with respect to B j if it is the first row in a slice and its leader belongs to B t with t ? j. Note that, all elements must be in either regular rows or special rows with respect to B j . At this point, we make a key observation. Observation 5.1. With respect to every bucket B j , there exist m regular rows and at most m special rows in R k;k . Proof. The number of regular rows follows directly from the definition of bucket B j in (2). The claim concerning the number of special rows follows from the assumed sortedness of the m n slices of size ( m implies that each slice of size ( m may contain at most one special row with respect to any bucket. In order to process each of the ( m buckets individually, we partition the mesh R k;k into submeshes T each of size ( m . Specifically, T 1 contains the leftmost m columns of R k;k , T 2 contains the next m columns, and so on. Each submesh is dedicated to bucket B j , in order to accumulate and process the elements belonging to that bucket, as we describe next. In O( m time we replicate the contents of T 1 in every submesh T Next, we broadcast in each submesh T j the values a (j \Gamma1) m and a j m that are used in (2) to define bucket B j . As a result, all the elements that belong to B j mark themselves. All the unmarked elements change their value to +1. At this point, it is useful to view the mesh R k;k as consisting of submeshes Q each of size m \Theta m . It is easy to see that processor P (r; s) is in Q l;j if n . The objective now becomes to move all the elements in T j belonging to bucket B j to the submesh Q j;j . To see how this is done, let q k (= a v ), be the leader of a regular row with respect to bucket B j . The rank r of this row is taken to be r 1. Now, in the order of their ranks, the regular rows with respect to are moved to the row in Q j;j corresponding to their rank. It is easy to confirm that all the regular rows with respect to B j can be moved into the submesh Q j;j in O( m time. Now consider a row u of T j that is special with respect to bucket B j . Row u is assigned the rank s= . Note that no two special rows can have the same rank. In the order of their ranks, special rows are moved to the rows of Q j;j corresponding to their ranks. As the number of special rows is at most m , the time taken to move all the special rows to Q j;j is O( m Notice that as a result of the previous data movement operations, each processor in Q j;j holds at most two elements: one from a regular row with respect to B j and one from a special row. Next, we sort the elements in each submesh Q j;j in overlaid row-major order. In case the number of elements in Q j;j does not exceed m 2 after sorting the elements can be placed one per processor. If the number of elements exceeds m 2 n , the first m 2 n of them are said to belong to generation-1 and the remaining elements are said to belong to generation- 2. The elements belonging to generation-1 are stored one per processor in row-major order, overlaid with those from generation-2, also in row-major order. This task is performed as follows. Using one of the optimal sorting algorithm for meshes [32, 45], sort the elements in regular rows in Q j;j in O( m repeat the same for the elements in special rows. Merging the two sorted sequences thus obtained can be accomplished in another O( m time. Now, in each submesh Q j;j , all the elements know their ranks in bucket B j . Our next goal is to compute the final rank of each of the elements in R k;k . Before we give the details of this operation, we let S 1 be the sorted slices of size ( m be the largest element in bucket B j . In parallel, using simple data movement, each m j is broadcast to all the processors in T j in O( m time. Next, we determine the rank of m j in each of the S l 's as follows: in every S l we identify the smallest element (if any) strictly larger than m j . Clearly, this can be done in at most O( m time, since every processor only has to compare m j with the element it holds and with the element held by its predecessor. Now the rank of m j among the elements in R k;k is obtained by simply adding up the ranks of m j in all the S l 's. Once these ranks are known, in at most O( m time they are broadcast to the first row of Q j;j , where their sum is computed in O(log m knows its rank in R k;k , every element in bucket B j finds, in O(1) time, its rank in R k;k by using its rank in its own bucket, the size of the bucket, and the rank of m j . Consequently, we have proved the following result. Lemma 5.2. The rank in R k;k of every element in every bucket can be determined in O( m time. Finally, we need to move all the elements to the leftmost m columns of R k;k in row-major order. In O(1) time, each element determines its final position from its rank r as follows. The row number x is given by d r e and the column number y by In every submesh T j , each element belonging to generation-1 is moved to the row x to which it belongs in sorted row-major order by broadcasting the m rows of Q j;j , one at a time. This takes O( m time. Notice that at this point every row of R k;k contains at most m elements. Knowing the columns to which they belong, in another m time all the elements can be broadcast to their positions along the row buses. This is repeated for the generation-2 elements. In parallel, every diagonal submesh R k;k moves back its data into the leftmost m columns of submesh R k;1 . Thus, in O( m time, all the elements are moved to the leftmost m columns of R. Now R contains slices of size ( m each sorted in row-major order. To summarize our findings we state the following result. Lemma 5.3. The task of merging m consecutive sorted slices of size ( m slice of size ( m can be performed in O( m time. be the worst-case complexity of the task of sorting a slice of size ( m It is easy to confirm that the recurrence describing the behavior of T (i The algorithm terminates at the end of t iterations, when Now, by dividing (1) throughout by by raising to the (t 1)-th power we obtain By combining (3) and (4), we obtain In turn, (5) implies that Thus, the total running time of our algorithm is given by which is obtained by solving the above recurrence. Since ffl is a constant, we have proved the following result. Theorem 5.4. For every choice of a constant set of m, n 1 elements stored in the leftmost d m e columns of a mesh with multiple broadcasting of size can be sorted in \Theta( m time. This is both time- and VLSI-optimal. 6 Conclusions and Open Problems The mesh-connected computer architecture has emerged as one of the most natural choices for solving a large number of computational tasks in image processing, computational geom- etry, and computer vision. Its regular structure and simple interconnection topology makes the mesh particularly well suited for VLSI implementation. However, due to its large communication diameter, the mesh tends to be slow when it comes to handling data transfer operations over long distances. In an attempt to overcome this problem, mesh-connected computers have been augmented by the addition of various types of bus systems. Among these, the mesh with multiple broadcasting (MMB) is of a particular interest being commercially available, being the underlying architecture of the DAP family of multiprocessors. The main contribution of this paper is to present the first known adaptive time- and VLSI-optimal sorting algorithm for the MMB. Specifically, we have shown that once we fix a constant the task of sorting m elements, n 1 pretiled in the leftmost m columns of an MMB of size can be performed in O( m time. This is both time- and VLSI-optimal. A number of problems remain open. First, it would be of interest to see whether the bucketing technique used in this paper can be applied to the problem of selection. To this day, no time-optimal selection algorithms for meshes with multiple broadcasting are known. Also, it is not known whether the technique used in this paper can be extended to meshes enhanced by the addition of k global buses [1, 12]. Further, we would like to completely resolve these issues concerning optimal sorting over the entire range that the results of Lin and others [24] show that for m near n,\Omega\Gamma378 n) is the time lower bound for sorting in this architecture. Their results imply that a sorting algorithm cannot be VLSI-optimal for m near Quite recently, Lin et al. [24] proposed a novel VLSI architecture for digital geometry - the Mesh with Hybrid Buses (MHB) obtained by enhancing the MMB with precharged 1-bit row and column buses. It would be interesting to see whether the techniques used in this paper extend to the MHB. This promises to be an exciting area for future work. Acknowledgement : The authors wish to thank the anonymous referees for their constructive comments and suggestions that led to a more lucid presentation. We are also indebted to Professor Ibarra for his timely and professional way of handling our submission --R Optimal bounds for finding maximum on array of processors with k global buses Parallel Computation: Models and Methods Square meshes are not always optimal Design of massively parallel processor STARAN parallel processor system hardware Square meshes are not optimal for convex hull computation A fast selection algorithm on meshes with multiple broadcasting Convexity problems on meshes with multiple broadcasting The MasPar MP-1 architecture Designing efficient parallel algorithms on mesh connected computers with multiple broadcasting Efficient median finding and its application to two-variable linear programming on mesh-connected computers with multiple broadcasting Prefix computations on a generalized mesh-connected computer with multiple buses Array processor with multiple broadcast- ing Image computations on meshes with multiple broadcast A multiway merge sorting network IEEE Transactions on Computers Parallel Processing Letters The mesh with hybrid buses: an efficient parallel architecture for digital geometry IEEE Transactions on Parallel and Distributed Systems Computer Vision Optimal sorting algorithms on bus-connected processor arrays Methods for realizing a priority bus system A Guided Tour of Computer Vision Bitonic sort on a mesh-connected parallel computer Finding connected components and connected ones on a mesh-connected parallel computer Data broadcasting in SIMD computers Optimal convex hull algorithms on enhanced meshes A. new deterministic sampling scheme The AMT DAP 500 The Massively Parallel Processor Parallel Computing: Theory and Practice Fractal graphics and image compression on a SIMD processor Constant time BSR solutions to parenthesis matching The VLSI complexity of sorting Sorting on a mesh-connected parallel computer Foundations of Vision Algorithms for sorting arbitrary input using a fixed-size parallel sorting device --TR

parallel algorithms;VLSI-optimality;lower bounds;sorting;meshes with multiple broadcasting;time-optimality

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Theoretical Analysis for Communication-Induced Checkpointing Protocols with Rollback-Dependency Trackability.

AbstractRollback-Dependency Trackability (RDT) is a property that states that all rollback dependencies between local checkpoints are on-line trackable by using a transitive dependency vector. In this paper, we address three fundamental issues in the design of communication-induced checkpointing protocols that ensure RDT. First, we prove that the following intuition commonly assumed in the literature is in fact false: If a protocol forces a checkpoint only at a stronger condition, then it must take, at most, as many forced checkpoints as a protocol based on a weaker condition. This result implies that the common approach of sharpening the checkpoint-inducing condition by piggybacking more control information on each message may not always yield a more efficient protocol. Next, we prove that there is no optimal on-line RDT protocol that takes fewer forced checkpoints than any other RDT protocol for all possible communication patterns. Finally, since comparing checkpoint-inducing conditions is not sufficient for comparing protocol performance, we present some formal techniques for comparing the performance of several existing RDT protocols.

Introduction A distributed computation consists of a finite set of processes connected by a communication network, that communicate and synchronize by exchanging messages through the network. A local checkpoint is a snapshot of the local state of a process, saved on nonvolatile storage to survive process failures. It can be reloaded into volatile memory in case of a failure to reduce the amount of lost work. When a process has to record such a local state, we say that this process takes a (local) checkpoint. With each distributed computation is thus associated a checkpoint and communication pattern, defined from the set of messages and local checkpoints. A global checkpoint [1] is a set of local checkpoints, one from each process, and a global checkpoint M is consistent if no message is sent after a checkpoint of M and received before another checkpoint of M [2]. The computation of consistent global checkpoints is an important work when one is interested in designing or implementing systems that have to ensure dependability of the applications they run. Many protocols have been proposed to select local checkpoints in order to form consistent global checkpoints (see the survey [3]). Remark that if local checkpoints are taken independently there is a risk that no consistent global checkpoint can ever be formed from them (this is the well-known unbounded domino effect , that can occur during rollback-recovery [4]). To avoid the domino effect, a kind of coordination in the determination of local checkpoints is required. In [2, 5], the coordination is achieved at the price of synchronization by means of additional control messages. Another approach, namely, communication-induced checkpointing [6], achieves coordination by piggybacking control information on application messages. In that case, processes select local checkpoints independently (called basic checkpoints) and the protocol requires them to take additional local checkpoints (called forced checkpoints) in order to ensure the progression of a consistent recovery line. Forced checkpoints are taken according to certain condition tested each time a message is received, on the basis of control information piggybacked on messages. Generally, the fact that two local checkpoints be not causally related is a necessary but not sufficient condition for them to belong to the same consistent global checkpoint [7]. They can have "hidden" dependencies that make them impossible to belong to the same consistent global checkpoint. These dependencies are characterized by the fact that they cannot be tracked with transitive dependency vectors. To solve this problem, Wang has defined the Rollback-Dependency Trackability (RDT) property [8]. A checkpoint and communication pattern satisfies this property if all dependencies between local checkpoints can be on-line trackable (i.e. trackable by a simple use of the transitive dependency vector). RDT has two noteworthy properties: (1) It ensures that any set of local checkpoints that are not pairwise causally related can be extended to form a consistent global checkpoint; (2) It enjoys efficient calculations of the minimum and the maximum consistent global checkpoints that contain a given set of local checkpoints. As a consequence, the RDT property has applications in a large family of dependability problems such as: software error recovery [9], deadlock recovery [10], mobile computing [11], distributed debugging [12], etc. Moreover, when combined with an appropriate message logging protocol [13], RDT allows to solve some dependability problems posed by nondeterministic computations as if these computations were piecewise deterministic [8]. Since the RDT property has a wide range of applications on many problems, it becomes an important pragmatic issue to design an efficient communication-induced checkpointing protocols satisfying the RDT property. The number of forced checkpoints and the size of piggybacked control information are dominant factors on the price to be paid. Hence the main question in this context is: how to design an efficient RDT protocol with less number of forced checkpoints and smaller size of piggybacked control information? Is the common intuition in the literature [8, 14], "If a protocol forces a checkpoint at a weaker condition then it must force at least as many checkpoints as a protocol that does at a stronger condition.", necessarily true (note that the stronger condition is a subset of the weaker condition)? Is there actually a tradeoff between the number of forced checkpoints and the size of piggybacked control information [15]? In this paper, we give a theoretical analysis for these problems. First, some counterexamples against the previous two statements are enumerated. We then demonstrate that there is no optimal on-line protocol in terms of the number of forced checkpoints. Since the common intuition is proved to be invalid, some interesting techniques of comparing useful protocols are also proposed. These techniques can be used to compare many existing RDT protocols in the literature. Remark that our results are not only from a theoretical point of view, but also from a practical one, when considering the task of designing efficient protocols with the RDT property. This paper is structured in five main sections. Section 2 defines the computational model and introduces definitions and elements of the Rollback-Dependency Trackability theory. In Section 3, we discuss some "impossibility " problems. Then several techniques of comparing useful protocols are addressed in the next section. Section 5 depicts a hierarchy graph for comparing a family of RDT protocols to marshal the discussions in the context. Finally, we conclude the paper in Section 6. 2. Preliminaries 2.1 Checkpoint and Communication Patterns A distributed computation consists of a finite set P of n processes fP that communicate and synchronize only by exchanging messages. We assume that each ordered pair of processes is connected by an asynchronous, reliable, directed logical channel whose transmission delays are unpredictable but finite. Each process runs on a processor; processors do not share a common memory; there is no bound for their relative speeds and they fail according to the fail-stop model. A process can execute internal, send and receive statements. An internal statement does not involve communication. When P i executes the statement "send(m) to puts the message m into the channel from P i to P j . When P i executes the statement "receive(m)", it is blocked until at least one message directed to P i has arrived; then a message is withdrawn from one of its input channels and received by P i . Executions of internal, send and receive statements are modeled by internal, sending and receiving events. Processes of a distributed computation are sequential , in other words, each process P i produces a sequence of events. This sequence can be finite or infinite. All the events produced by a distributed computation can be modeled as a partially ordered set with the well-known !", defined as follows [16]. Definition 1: The relation " hb !" on the set of events satisfies the following three condition: (1) If a and b are events in the same process and a comes before b, then a hb b. (2) If a is the event send(m) and b is the event receive(m), then a hb b. (3) If a hb a hb With each distributed computation is associated a checkpoint and communication pat- tern, which is composed of a distributed computation H and the set of local checkpoints !!!!!!!!!!!!! I k,1 I k,2 Figure 1: A checkpoint and communication pattern ccpat. defined on H. Figure 1 shows an example checkpoint and communication pattern ccpat. C i;x represents the xth checkpoint of process and the sequence of events occurring at P i between C called a checkpoint interval and is denoted by I i;x ; where i is called the process id and x the index of this checkpoint or this checkpoint interval. We assume each process P i starts its execution with an initial checkpoint C i;0 . 2.2 Rollback-Dependency Trackability First we will briefly introduce the concepts of Z-path [7] and causal doubling of a Z-path [15], and then the concept of Rollback-Dependency Trackability [8]. For more details on these subjects, please consult those papers previously cited. Definition 2: A Z-path is a sequence of messages [m each t. To the best of our knowledge this notion has been introduced for the first time by Netzer and Xu in [7] under the name zig-zag path. If a Z-path [m is such that we say that this Z-path is from I i;x to I j;y . For example, in pattern ccpat depicted in Figure 1, both the paths [m are Z-path from I k;2 to I i;2 . However the path [m is not a Z-path. In the rest of this paper, we use the following notation. In a Z-path -, the first (last) message will be denoted -:first (-:last). Let - and - be two Z-paths whose concatenation is also a Z-path. This concatenation will be represented as - \Delta -. We now introduce the notion of a causal Z-path. A Z-path is causal if the receiving event of each message except the last one precedes the sending event of the next message in the sequence. A Z-path is non-causal if it is not causal. A Z-path with only one message is trivially causal. For simplicity, a casual Z-path will also be called a causal path. Definition 3: A Z-path from I i;x to I j;y is causally doubled if or if there exists a causal path - from I i;x 0 to I j;y 0 where x - x 0 and y 0 - y. From the previous definition, every causal path is obviously causally doubled by itself. As an example, the Z-path [m in pattern ccpat of Figure 1 is non-causal, and is causally doubled by the causal path [m The following concept, Rollback-Dependency Trackability , was introduced by Wang in [8]. It is defined differently but equivalently as below [15]. Definition 4: A checkpoint and communication pattern ccpat satisfies the RDT property if and only if all non-causal Z-paths are causally doubled in ccpat. In other words, a checkpoint and communication pattern satisfies this property in the sense that all dependencies between local checkpoints need to be on-line trackable since dependencies can be passed over only along causal paths. 2.3 PCM-paths For a given checkpoint and communication pattern ccpat, it is not necessary to check that every non-causal Z-path is causally doubled to ensure that ccpat satisfies the RDT property. Namely, we can only consider some certain embedded subsets of non-causal Z-paths [15]. An important subset, PCM-paths, is now introduced. Definition 5: A causal path - from I i;x to I j;y is prime if every causal path - from I i;x 0 to I j;y 0 with x - x 0 and y 0 - y satisfies that receive(-:last) hb receive(-:last). Intuitively, a prime path from I i;x to I j;y is the first one including the existence of interval I i;x (i.e. the new dependency on which P j 's current state transitively depends with the causal past of P j 's current state. A PCM-path - \Delta m is a Z-path that is the concatenation of a causal path - and a single message m, where - is prime and send(m) hb receive(-:last). In pattern ccpat shown in Figure 1, the path [m 5 ] is prime but [m 3 ] is not prime. And the path [m is a PCM-path. The following theorem is a direct consequence of [15]. Theorem 1: A checkpoint and communication pattern ccpat satisfies the RDT property if and only if all PCM-paths are causally doubled in ccpat. The idea behind this theorem is that according to the definitions of RDT, since all dependencies between local checkpoints must be on-line trackable, any new dependency has to be passed over along a Z-path to its end. Due to all new dependencies included by prime causal paths firstly, that all PCM-paths are causally doubled is then necessary and sufficient to ensure the RDT property. Note that we can exploit the following transitive dependency tracking mechanism proposed in the literature [11, 17, 18] to detect the existence of a prime path: In a system with processes, each process P i maintains the size-n transitive dependency vector (TDV), that can represent the current interval index (or equivalently the checkpoint index of the next checkpoint) of P i and record the highest index of intervals of any other process P j on which current state transitively depends. TDV is piggybacked on application messages sent, and upon the receipt of messages, processes can decide by evaluating this vector if a prime path is encountered [8]. According to Theorem 1, in order to ensure the RDT property, any PCM-path which is not causally doubled needs to be broken by a forced checkpoint. For a PCM-path - \Delta m whose breakpoint is I i;x , on the receipt of -:last, process P i has to distinguish if - \Delta m is causally doubled, only through the information carried on the message. So this PCM-path has to be not only causally doubled but also visibly doubled , defined as follows [15], in order not to be broken. Definition visibly doubled if and only if it is causally doubled by a causal path - 0 with receive(- 0 :last) hb send(-:last). Figure 2 shows an example visibility of doubling. Intuitively a causal doubling of a PCM- path is visible at a process on the receipt of message -:last, if the path - 0 that causally doubles belongs to the causal past of -:last. Note that from the definition a causally doubled PCM-path is not necessarily visibly doubled, but a non-causally doubled one must be non-visibly doubled. Based on the foregoing discussion and Theorem 1, we can deduce a characterization of the RDT with respect to protocols based on the entire-causal history [15]. Corollary 1: A checkpoint and communication pattern produced by a protocol based on the entire-causal history satisfies the RDT property if and only if all PCM-paths are visibly doubled. -.last Figure 2: Visibility of doubling. If a PCM-path is from an interval I i;x to another interval I i;x 0 of the same process P i , we call this PCM-path a PCM-cycle. If x PCM-cycle can not be causally doubled, and is called a non-doubled PCM-cycle [15]. For example, the path [m Figure 1 is a PCM-cycle and is non-doubled. In the remainder of the paper, for the sake of clarity, we only refer to a PCM-path from a process to another different one as a PCM-path, and on the contrary, the PCM-path from a process to the same one as a PCM-cycle. PCM-paths and PCM-cycles are all called PCM-conditions. 3. Impossibility Problems In this section, we discuss some "impossibility " problems. First we disprove the truthfulness of the following common intuition in the literature [8, 14]: if a protocol forces a checkpoint at a weaker condition then it must force at least as many checkpoints as a protocol that does at a stronger condition. In other words, even though conditions involved in two different protocols have inclusive and subordinative relationship, it may be impossible to compare these two protocols in terms of the number of forced checkpoints. The motive of this problem is that since any forced checkpoint will change the given checkpoint and communication pattern and consequently affect the later condition testings, as soon as two protocols differ in their decision to force a checkpoint, the two resulting checkpoint and communication patterns are not the same and possibly strongly "diverge" in the future, and thus it is perhaps impossible to compare these two protocols. We also overthrow the concept that there is a tradeoff between the number of forced checkpoints and the size of piggybacked control information for RDT protocols. In the last subsection, it was given a proof to demonstrate another impossibility problem that there is no optimal on-line protocol that ensures the RDT property. This scenario is quite common in the area of on-line algorithms due to no knowledge of future information. 3.1 Common Intuitions Not Necessarily True Two counterexamples are enumerated against those specious statements mentioned previ- ously. Their results show the fact that it definitely needs a formal proof for comparison of two different protocols. Therefore we will propose some techniques of comparing useful protocols in the next section. Counterexample 1: CPn is a protocol that breaks all PCM-paths and every PCM-cycle ! send(-:first), and CPm is a protocol that breaks all PCM-paths and every CM-cycle - \Delta m (a CM-condition is the concatenation of a causal path - which is not necessarily prime and a single message m where send(m) hb receive(m) hb ! send(-:first). It can be easily verified that both CPn and CPm break all non-causally doubled PCM-paths and non-doubled PCM-cycles. Hence these two protocols are RDT protocols by Theorem 1. And they only need to piggyback TDV on application messages (note that we apply the consequence of [15] to evaluate the size of piggybacked control information in this paper and please refer to that paper for more details). Obviously, CPm forces a checkpoint at a weaker condition than CPn's. As the result shown in Figure 3 (a), CPn must take two forced checkpoints (the diamond box) to make the considered checkpoint and communication pattern satisfy the RDT property. However, CPm needs only one forced checkpoint to make the same pattern ensure RDT, depicted in Figure 3 (b). This counterexample shows that CPm forces fewer checkpoints than CPn in the given checkpoint and communication pattern, and thus disproves the common intuition. Besides overthrowing the common intuition, the following counterexample also demonstrates that there is not necessarily a tradeoff between the number of forced checkpoints and the size of piggybacked control information. Counterexample 2: Let No-Non-Visibly-Doubled-PCM (NNVD-PCM) a protocol that breaks all non-visibly doubled PCM-paths and non-doubled PCM-cycles, and No- PCM-Path a protocol that breaks all PCM-path and non-doubled PCM-cycles. Similarly, since both protocols break all non-visibly doubled PCM-paths and non-doubled PCM-cycles, initial checkpoint forced checkpoint basic checkpoint (a) Figure 3: The scenario of Counterexample 1 (a) CPn (b) CPm. according to Corollary 1, they are also RDT protocols. Moreover, since NNVD-PCM has to decide whether a PCM-path is visibly doubled, it needs more piggybacked control information than No-PCM-Path's. NNVD-PCM takes one more forced checkpoint (see Figure 4 (a)) than No-PCM-Path does (see Figure 4 (b)). Hence this counterexample also shows that the protocol piggybacking less control information (No-PCM-Path) outperforms the one piggybacking more control information (NNVD-PCM) in some checkpoint and communication patterns in terms of the number of forced checkpoints. The idea behind those counterexamples is that the forced checkpoint taken by the protocol based on the stronger condition will make the CM-path at the rightmost part of the considered checkpoint and communication pattern become a non-causally doubled PCM-path, and thus another checkpoint is necessary to be forced to break this PCM-path. However, the forced checkpoint taken by the protocol on the weaker condition will not give rise to such a scenario. initial checkpoint forced checkpoint basic checkpoint (a) Figure 4: The scenario of Counterexample 2 (a) NNVD-PCM (b) No-PCM-Path. 3.2 No Optimal On-line Protocol We take two categories of on-line protocols into consideration for this problem. These two categories are on-line protocols based on the entire causal history and the transitive dependency tracking (i.e. only piggybacking TDV on a message as control information), respec- tively. They are all shown to have no optimal protocol through the following descriptions. Given the checkpoint and communication pattern in Counterexample 1, redrawn in Figure 5 and denoted ccpat a , we directly have the following lemma since the PCM-cycle [m is non-doubled. Lemma 1: Process P 3 needs to force at least one checkpoint between point a and b to make ccpat a satisfy RDT for all entire-causal, and TDV on-line RDT protocols. Lemma 2: If a forced checkpoint is taken between point c and b in Process P 3 , Process P 2 must take another forced checkpoint to satisfy RDT for all entire-causal, and TDV on-line RDT protocols. ccpat c ccpat a initial checkpoint m 4 a c b Figure 5: The checkpoint and communication patterns ccpat a and ccpat c . forced checkpoint is taken between point c and b in Process P 3 , the Z-path becomes a non-causally doubled PCM-path since m 3 turns out prime. Hence P 2 has to force another checkpoint to break this PCM-path to satisfy RDT for all entire-causal, and TDV on-line RDT protocols, as the scenario shown in Figure 3 (a). Q.E.D. We now consider the following theorem. Theorem 2: There is no optimal on-line protocol based on the entire causal history in terms of the number of forced checkpoint. protocol is optimal if and only if given any checkpoint and communication pat- tern, no other protocol has less number of forced checkpoints than it. See the checkpoint and communication pattern ccpat a depicted in Figure 5. Since the protocol CPm in Counterexample needs only one forced checkpoint to make ccpat a satisfy RDT, protocols that force any checkpoint between point c and b (by Lemma 2 these protocols must force two checkpoints), and protocols that take more than one checkpoint between point a and c cannot be optimal. So an optimal protocol, if any, must take exactly one forced checkpoint between point a and c according to Lemma 1. However, such an on-line protocol cannot be optimal because it has to force one checkpoint in the cut pattern ccpat c , shown as the left region of the dotted line in Figure 5, due to the same causality with that in ccpat a at the same position. But the protocol CPn in Counterexample 1 takes zero forced checkpoint in ccpat c . Therefore, there is no optimal on-line protocol based on the entire causal history in terms of the number of forced checkpoints. Q.E.D. As mentioned earlier in Counterexample 1, both the protocols CPm and CPn only piggyback the transitive dependency vector. Hence we can have the corollary below with the similar proof of the previous theorem. Corollary 2: There is no optimal on-line protocol based on the transitive dependency tracking (TDV) in terms of the number of forced checkpoints. 4. Techniques of Comparison 4.1 FDAS vs Other Protocols Wang proposed the Fixed-Dependecy-After-Send (FDAS) checkpointing protocol in [8], that breaks all PCM-conditions. In this subsection, two assertions that FDAS outperforms protocols which force a checkpoint at weaker conditions than FDAS's, and that protocols which force a checkpoint at stronger conditions are better than FDAS, both in terms of the number of forced checkpoints, are demonstrated. First we prove that the former assertion is true. Let C f denotes the condition on which FDAS is based (i.e. breaking all PCM- conditions), and CPw denotes a protocol which takes a forced checkpoint at a weaker condition than FDAS's, with its based condition denoted Cw . Obviously C f is a subset of Cw and we represent this relation as C f ) Cw . Let ccpat f and ccpat w represent the checkpoint and communication patterns produced by the protocols FDAS and CPw respectively. Since adding forced checkpoints cannot make any PC-path (i.e. prime causal path) in the original checkpoint and communication pattern become a non-PC path, we directly have the following lemma. Lemma 3: For any PC-path in the original checkpoint and communication pattern, it is still a PC-path in the checkpoint and communication pattern produced by any protocol. Now we define an extra PC-path as a non-PC path originally but becoming a PC-path due to the forced checkpoint taken by a protocol. We then have the following two lemmas. Lemma 4: For any extra PC-path in ccpat f , it is also an extra PC-path in ccpat w . Proof: Assume there exists an extra PC-path - in ccpat f which is not an extra PC-path in ccpat w . An extra PC-path is produced only when a forced checkpoint is taken before it. Thus there exists a PC-path condition - 1 before send(-:first) in ccpat f such that FDAS forced the process to take a checkpoint to break the C f condition formed by - 1 , and this checkpoint made - become an extra PC-path. We then say that - is produced by - 1 . Obviously, - 1 is in the causal past of -. Similarly, - 1 is a PC-path either because itself is originally a PC-path Figure The observation process in Lemma 4. or because it is an extra PC-path produced by another PC-path - 2 , which is in the causal past of - 1 and thus also in the causal past of -. By repeatedly the foregoing observation, and since messages in the causal past of - is finite, eventually we can obtain an original PC-path - n . The previous progress is shown in Figure 6. By Lemma 3, - n is also a PC-path in ccpat w . Hence CPw must force the process to take a checkpoint (not necessarily the same checkpoint with the one in ccpat f ) between receive(- n :last) and its nearest previous message-sending event to avoid - n to form a C f condition that CPw also needs to break since C f ) Cw . The forced checkpoint makes - n\Gamma1 in ccpat w become an extra PC-path since - in ccpat f is also an extra PC-path and there is no other message-sending event between receive(- n :last) and its nearest previous message-sending event (namely there cannot exist a causal path that prevents - n\Gamma1 from becoming prime). For the same reason, - n\Gamma2 also becomes a PC-path in ccpat w , and finally - also becomes an extra PC-path in ccpat w . This leads to a contradiction. Q.E.D. Lemma 5: FDAS can never force two consecutive checkpoints between two consecutive Condition C w (if any) PC-path PC-path Figure 7: The scenario of Lemma 5. checkpoints forced by CPw. Proof: We prove this lemma by showing that CPw must force at least one checkpoint between any two consecutive forced checkpoints taken by FDAS. See the scenario shown in Figure 7, there are two consecutive checkpoints forced by FDAS, so there exist two continuous conditions. For the PC-path of the latter C f condition in ccpat f , by Lemma 3 and Lemma 4, it is also a PC-path in ccpat w . Therefore CPw has to force one checkpoint between this PC-path and its nearest previous message-sending event to prevent a C f condition from being formed, shown as the hollow diamond in Figure 7. This checkpoint is obviously between the two consecutive checkpoints forced by FDAS. Q.E.D. As a consequence, we can derive the following "monotonicity " property. Theorem 3: CPw takes the n-th forced checkpoint no later than FDAS does, for all n. Proof: By induction, because the given checkpoint and communication pattern are exactly the same for CPw and FDAS before any forced checkpoint is taken, it is clear that CPw must force the first checkpoint no later than FDAS does for the reason that C f ) Cw . Now suppose CPw takes the k-th checkpoint no later than FDAS does, according Lemma 5, we have that CPw will take the 1)-th checkpoint no later than FDAS does. Q.E.D. Let #f ckpt(CP) denotes the number of forced checkpoints taken by the protocol CP. Applying the previous theorem, it is obvious that #f ckpt(CPw) - #f ckpt(FDAS). The Russell's algorithm [19] and the protocol presented in [20], that are named No-Receive- After-Send (NRAS) and Fixed-Dependency-Interval (FDI) by Wang in [8] respec- tively, are both RDT protocols. By their definitions, NRAS breaks all CM-paths and FDI forces a checkpoint whenever a PC-path is encountered. Hence they belong to the family of CPw. By Theorem 3, we know that FDAS is better than these two protocols in terms of the number of forced checkpoints. Here we begin to demonstrate that the latter assertion aforementioned in the very beginning is valid. Also, let CPs denote a protocol which takes a forced checkpoint at a stronger condition than FDAS's, and the condition it based is C s , obviously where C s ) C f . Let ccpat s represent the checkpoint and communication pattern produced by the protocol CPs. Similarly we consider the following lemma. Lemma For any extra PC-path in ccpat s , it is also an extra PC-path in ccpat f . Proof: Assume there exists an extra PC-path - in ccpat s , which is not an extra PC-path in ccpat f . Since CPs also only breaks some certain PCM-conditions, by the same observation with Lemma 4, we can obtain an original PC-path - n , that causes - to become an extra PC-path. By Lemma 3, - n is also a PC-path in ccpat f . Hence FDAS must force the process to take a checkpoint (not necessarily the same checkpoint with the one in ccpat s ) between receive(- n :last) and its nearest previous message-sending event to avoid - n to form a C f condition that FDAS needs to break. The forced checkpoint makes - n\Gamma1 in ccpat f become an extra PC-path since - n\Gamma1 in ccpat s is also an extra PC-path and there is no other message-sending event between receive(- n :last) and its nearest previous message-sending event. For the same reason, - n\Gamma2 also becomes a PC-path in ccpat f , and finally - also becomes an extra PC-path in ccpat f . This leads to a contradiction. Q.E.D. With the similar description of Lemma 5, the following lemma can be proved. Lemma 7: CPs can never force two consecutive checkpoints between two consecutive checkpoints forced by FDAS. Therefore we can obtain the corollary below in a straightforward way. Corollary 3: FDAS takes the n-th checkpoint no later than CPs does, for all n. And so, the assertion #f ckpt(FDAS) - #f ckpt(CPs) holds. The RDT protocol BHMR proposed in [14] breaks non-visibly doubled PCM-paths and some PCM-cycles (including all non-doubled PCM-cycles), and thus it belongs to the family of CPs. As a side effect of the foregoing corollary, we give a formal proof showing that BHMR outperforms FDAS in terms of the number of forced checkpoints, instead of the simulation results in 4.2 No-PCM-Cycle vs FDAS Another interesting result is the technique for comparing No-PCM-Cycle and FDAS. Applying Corollary 3, the protocol No-PCM-Cycle that breaks non-visibly doubled PCM- paths and all PCM-cycles outperforms FDAS because its based condition is stronger than FDAS's. However we find that No-PCM-Cycle is actually equivalent to the protocol FDAS, shown in the following theorem. Theorem 4: If all PCM-cycles and non-visibly doubled PCM-paths are broken in the check-point and communication pattern, any visibly doubled PCM-path is also broken. Proof: For a visibly doubled PCM-path - \Delta m shown in Figure 8 (a), since there must exist a prime path that visibly doubles - \Delta m, without loss of generality, - 1 can be assumed prime. From Figure 8 (a), we know that - 1 is in the causal past of -:last. Now we show by a case analysis that - \Delta m will be broken. (a). If - 1 \Delta (-:last) is prime, the PCM-cycle - 1 \Delta (-:last) \Delta m is broken, and thus the PCM-path (b). If - 1 \Delta (-:last) is not prime, there exists a causal path - 0 1 to process P i , which is necessarily before receive(- 1 :last) (otherwise - will turn out non-prime), as shown in Figure 8 (a). Without lost of generality, we assume - 0 1 to be the nearest causal path to P i before 1 :first) is a PCM-condition. If this PCM-condition is broken by the theorem assumption, the forced checkpoint will make - 1 \Delta (-:last) become prime since is the first causal path to P i after receive(- 1 :last), and consequently the will be broken. If not, the PCM-path 1 :first) is a visibly doubled PCM-path, shown as Figure 8 (b). Clearly, - 2 is in the causal past of - 1 :last and thus in the causal past of - 1 (also in the causal past of -:last). By repeatedly applying the foregoing observation, and since messages in the causal past of -:last is finite, eventually we can obtain a PCM-condition - n \Delta (- 0 n :first), which is either a non-visibly doubled PCM-path or a PCM-cycle and therefore both need to be broken. The forced checkpoint will make become prime and the PCM-path - will be broken. With the same reason, - n\Gamma2 \Delta (- 0 n\Gamma2 :first) will be also broken and finally the PCM-path - \Delta m will also be broken. Q.E.D. According to the previous theorem, we have that the protocol No-PCM-Cycle in fact breaks all PCM-conditions, so it is equivalent to FDAS. This result shows that we can '.first .last (a) (b) Figure 8: The scenario of Theorem 4. reduce the redundant size of piggybacked control information by adopting FDAS instead of No-PCM-Cycle because No-PCM-Cycle needs extra information to distinguish a visibly doubled PCM-path. 4.3 PCM vs PESCM In [15], Baldoni et. al. proposed a more constrained characterization of the RDT property, PESCM , for designing protocols. A PESCM-condition is composed of a PCM-condition - \Delta m such that - is elementary and simple besides being prime. The interest of this subsection lies in the fact that the existence of a PCM-condition implies the existence of a PESCM- condition at the same position. It consequently becomes unnecessary to take such a stronger condition with more piggybacked control information into consideration for some protocols. First, we introduce the definitions of the terms "elementary" and "simple" [15]. Definition 7: A Z-path - is elementary if its traversal sequence P is the sequence of processes traversed by -, has no repetition. Definition 8: A causal path simple if the two events receive(m i ) and send(m i+1 ) occur in the same interval, 8i (1 Namely, an elementary Z-path only traverses a process once, and a simple causal path does not include local checkpoints. For instance, in the checkpoint and communication pattern shown in Figure 1, the path [m neither elementary nor simple because it traverses process P j twice and the local checkpoint C k;1 is included. But the path [m is both elementary and simple. By Definition 7, we have that every causal path contains an elementary causal path. The elementary causal path contained by a prime causal path - has the same starting interval and ending point with - and thus is also prime. A PECM-condition is defined as a PCM- condition with the property that - is elementary in addition to being prime. Then we directly have the following theorem and corollary. Theorem 5: The existence of a PCM-condition implies the existence of a PECM-condition at the same position. Corollary 4: The existence of a non-doubled PCM-condition implies the existence of a non-doubled PECM-condition at the same position. Next we begin to demonstrate the lemma below. Lemma 8: A PEC-path (prime and elementary causal path) contains a PESC-path (prime, elementary and simple causal path). Proof: Note that a non-simple causal path - can be written as each component - i is simple. We prove this lemma by showing that the last simple path contained by a PEC-path is prime (of course is elementary). As depicted in Figure 9, a PEC-path l is not prime, then there exists a prime path - from point x 0 to point y 0 , where point y 0 precedes point y. Since - 1 is a causal path with receive(-:last) hb turns out non-prime. This leads to a contradiction. Q.E.D. According to Theorem 5, Corollary 4 and the previous lemma, it can be easily verified that the following theorem and corollary are true. Theorem 6: The existence of a PCM-condition implies the existence of a PESCM-condition at the same position. Corollary 5: The existence of a non-doubled PCM-condition implies the existence of a non-doubled PESCM-condition at the same position. The idea underlying the previous results is that the protocol No-PESCM that breaks all PESCM-conditions and the protocol No-Non-Visibly-Doubled-PESCM that breaks non-visibly doubled PESCM-paths and non-doubled PESCM-cycles presented in [15] are exactly the same with FDAS and NNVD-PCM respectively, however with more piggybacked control information for the sake of necessity to distinguish simple paths. Intuitively, we have to break all non-doubled PCM-conditions in order to satisfy RDT by Theorem 1. Why can Figure 9: The PEC-path we achieve this goal only by breaking non-doubled PESCM-conditions in a scenario that there is not necessarily a PESCM-condition before a PCM-condition such that breaking the former can eliminate the latter? The reason is because the previous results hold. 5. A Family of RDT Protocols Figure depicts a hierarchy graph of comparing a family of communication-induced check-pointing protocols satisfying the RDT property. A plain arrow from a protocol CP1 to another protocol CP2 indicates that #f ckpt(CP1) - #f ckpt(CP2) and a dotted arrow indicates that the piggybacked control information of CP1 is less than that of CP2. The line with two arrows means "equivalent " and the line with a mark "X" on it means "incomparable". For the protocol CBR (Checkpoint-Befoe-Receive) [8] at the bottom of Figure 10, a checkpoint is placed before every message-receiving event. It can be easily verified that both FDI and NRAS force fewer number of checkpoints than CBR. Figure 10 marshals the discussions in the previous sections. Note that this family includes many existing RDT protocols in the literature. Therefore the result is helpful for a wide range of practical applications. No-PCM-Path No-PCM-Circle FDAS NRAS FDI CBR Figure 10: Comparing a family of RDT protocols. 6. Conclusions This paper has provided a theoretical analysis for RDT protocols. In the context, some "impossibility " problems are addressed. First, we have shown that the common intuitions in the literature are not convincing and it definitely needs to a formal proof to demonstrate that one protocol is better than another one. Through rigorous demonstrations, we usually found that it is not necessarily worthy to adopt protocols based on stronger conditions with more piggybacked control information. We also proved that there is no optimal on-line protocol that ensures the RDT property. This scenario is quite common in the area of on-line algorithms due to no knowledge of future information. Moreover, some techniques for comparing useful protocols have been proposed. We showed that these techniques can be exploited to compare many existing protocols in the literature. Hence our results provide guidelines for designing and evaluating efficient communication-induced checkpointing RDT protocols. Acknowledgements The authors wish to express their sincere thanks to Jean-Michel Helary (IRISA) and Michel Raynal (IRISA) whose comments helped improve the presentation of the paper. We would like also to thank Jeff Westbrook for his valuable discussions about on-line algorithms. Tsai and Kuo's work are supported by the National Science Council, Taiwan, ROC, under Grant NSC 87-2213-E-259-007. --R "Consistent global checkpoints based on direct dependency tracking," "Distributed snapshots: determining global states of distributed systems," "A survey of rollback-recovery protocols in message-passing systems," "System structure for software fault-tolerant," "Checkpointing and rollback-recovery for distributed systems," "Experimental evaluation of multiprocessor cache-based error recovery," "Necessary and sufficient conditions for consistent global snapshots," "Consistent global checkpoints that contain a given set of local check- points," "The maximum and minimum consistent global checkpoints and their applications," "Guaranteed deadlock recovery: deadlock resolution with rollback propagation," "Checkpointing distributed applications on mobile computers," "Causal distributed breakpoints," "When piecewise determinism is almost true," "A communication-induced checkpointing protocol that ensures rollback-dependency trackability," "Rollback-dependency trackability: an optimal characterization and its protocol," "Time, clocks and the ordering of events in a distributed system," "Efficient distributed recovery using message logging," "Optimistic recovery in distributed systems," "State restoration in systems of communicating processes," "Optimal checkpointing and local recording for domino-free rollback recovery," --TR --CTR Roberto Baldoni , Jean-Michel Hlary , Michel Raynal, Rollback-dependency trackability: visible characterizations, Proceedings of the eighteenth annual ACM symposium on Principles of distributed computing, p.33-42, May 04-06, 1999, Atlanta, Georgia, United States Jichiang Tsai, On Properties of RDT Communication-Induced Checkpointing Protocols, IEEE Transactions on Parallel and Distributed Systems, v.14 n.8, p.755-764, August B. Gupta , Z. Liu , Z. Liang, On designing direct dependency: based fast recovery algorithms for distributed systems, ACM SIGOPS Operating Systems Review, v.38 n.1, p.58-73, January 2004 B. Gupta , S. K. Banerjee, A Roll-Forward Recovery Scheme for Solving the Problem of Coasting Forward for Distributed Systems, ACM SIGOPS Operating Systems Review, v.35 n.3, p.55-66, July 1 2001 D. Manivannan , Mukesh Singhal, Quasi-Synchronous Checkpointing: Models, Characterization, and Classification, IEEE Transactions on Parallel and Distributed Systems, v.10 n.7, p.703-713, July 1999 J.-M. Hlary , A. Mostefaoui , R. H. B. Netzer , M. Raynal, Communication-based prevention of useless checkpoints in distributed computations, Distributed Computing, v.13 n.1, p.29-43, January 2000 Jun-Lin Lin , Margaret H. Dunham, A Low-Cost Checkpointing Technique for Distributed Databases, Distributed and Parallel Databases, v.10 n.3, p.241-268, December 2001

communication-induced protocols;checkpointing;rollback recovery;rollback-dependency trackability;on-line algorithms

291369

Diskless Checkpointing.

AbstractDiskless Checkpointing is a technique for checkpointing the state of a long-running computation on a distributed system without relying on stable storage. As such, it eliminates the performance bottleneck of traditional checkpointing on distributed systems. In this paper, we motivate diskless checkpointing and present the basic diskless checkpointing scheme along with several variants for improved performance. The performance of the basic scheme and its variants is evaluated on a high-performance network of workstations and compared to traditional disk-based checkpointing. We conclude that diskless checkpointing is a desirable alternative to disk-based checkpointing that can improve the performance of distributed applications in the face of failures.

Introduction Checkpointing is an important topic in fault-tolerant computing as the basis for rollback recovery. Suppose a user is executing a long-running computation and for some reason (hardware or software), the machine running the computation fails. In the absence of checkpointing, when the machine becomes functional, the user must start the program over, thus wasting all previous computation. Had the user stored periodic checkpoints of the program's state to stable storage, then he or she could instead restart the program from the most recent checkpoint. This is called rolling back to a stored checkpoint. For long-running computations, checkpointing allows users to limit the amount of lost computation in the event of a failure (or failures). There have been many programming environments intended for users with long-running computations that rely on checkpointing for fault-tolerance. For example, Condor [34], libckpt [25] and others [16, 30, 37] provide [email protected]. This material is based upon work supported by the National Science Foundation under grants CCR-9409496, MIP-9420653 and CDA-9529459, by the ORAU Junior Faculty Enhancement Award, and by DARPA under grant N00014-95-1-1144 and contract DABT63-94-C-0049. transparent checkpointing for uniprocessor programs, and checkpointers such as MIST [4], CoCheck [33] and others [2, 10, 18, 28, 32] provide checkpointing in parallel computing environments. All of the above systems store their checkpoints on stable storage (i.e. disk), since stable storage typically survives processor failures. However, since checkpoints can be large (up to hundreds of megabytes per processor), the act of storing them to disk becomes the main component that contributes to the overhead, or performance degradation, due to checkpointing. This is more marked in parallel and distributed systems where the number of processors often vastly outnumbers the number of disks. Several techniques have been devised and implemented to minimize this source of overhead, including incremental checkpointing [11, 38], checkpoint buffering with copy-on-write [9, 21], compression [20, 28] and memory exclusion [25]. However with all of these techniques, the performance of the stable storage medium is still the underlying cause of overhead. In this paper, we present diskless checkpointing. The goal of diskless checkpointing is to remove stable storage from checkpointing in parallel and distributed systems, and replace it with memory and processor redundancy. By eliminating stable storage, diskless checkpointing removes the main source of overhead in checkpointing. However, this does not come for free. The failure coverage of diskless checkpointing is less than checkpointing to stable storage, since none of the components in a diskless checkpointing system can survive a wholesale failure. Moreover, there is memory, processor and network overhead introduced by diskless checkpointing that is absent in standard disk-based schemes. The purpose of this paper is twofold. We first present basic schemes for diskless checkpointing and then performance optimizations to the basic schemes. Second, we assess the performance of diskless checkpointing on a network of Sparc-5 workstations as compared to standard disk-based checkpointing. As anticipated, diskless checkpointing induces less overhead on applications than disk-based checkpointing, enabling the user to checkpoint more frequently without a performance penalty. This lowers the application's expected running time in the presence of failures. Diskless checkpointing tolerates single processor failures, and in some cases multiple processor failures. How- ever, it does not tolerate wholesale failures (such as a power outage that knocks out all machines). Thus, an optimized fault-tolerant scheme would be a two-level scheme, as advocated by Vaidya [35], where diskless checkpoints are taken frequently, and standard, disk-based checkpoints are taken at a much larger interval. In this way, the more frequent case of one or two processors failing is handled swiftly, with low overhead, while the rarer case of a wholesale failure is handled as well, albeit with higher overhead and a longer rollback penalty. 2 Overview of Diskless Checkpointing Diskless checkpointing is based on coordinated checkpointing. With coordinated checkpointing, a collection of processors with disjoint memories coordinates to take a checkpoint of the global system state. This is called a "coordinated checkpoint". A coordinated checkpoint consists of checkpoints of each processor in the system plus a log of messages in transit at the time of checkpointing. Coordinated checkpointing is a well-studied topic in fault-tolerance. For a thorough discussion of coordinated checkpointing, the reader is directed to the survey paper by Elnozahy, Johnson and Wang [8]. With diskless checkpointing we assume that there is no message log to be stored (for example, the "Sync-and- stop" algorithm for coordinated checkpointing ensures that there is no message log [28]), or that the message log is contained within the checkpoints of individual processors. This reduces the problem of taking a coordinated checkpoint to saving the individual checkpoints of each processor in the system. Diskless checkpointing is composed of two parts - (1) checkpointing the state of each application processor in memory, and (2) encoding these in-memory checkpoints and storing the encodings in checkpointing processors. When a failure occurs, the system is recovered in the following manner. First, the non-failed application processors roll themselves back to their stored checkpoints in memory. Next, replacement processors are chosen to take the place of the failed processors. Finally, the replacement processors use the checkpointed states of the non-failed application processors plus the encodings in the checkpointing processors to calculate the checkpoints of the failed processors. Once these checkpoints are calculated, the replacement processors roll back, and the application continues from the checkpoint. Note that either spare processors or some of the checkpointing processors may be used as replacement processors. If checkpointing processors are used, then the system will continue with fewer (or no) checkpointing processors, thus reducing the fault-tolerance. However, when more processors become available, they may be employed as additional checkpointing processors. 2.1 Exact Problem Specification The user is executing a long-running application on a parallel or distributed computing environment composed of processors with disjoint memories that communicate by message-passing. The application executes on exactly n processors. With diskless checkpointing, an extra m processors are added to the system, and the n+m processors cooperate to take diskless checkpoints. As long as the number of processors in the system is at least n, and as long as failures occur within certain constraints, the application may proceed efficiently. As stated above, diskless checkpointing may be broken into two parts: application processors checkpointing themselves, and checkpoint processors encoding the application processors' checkpoints. Each is explained below, followed by issues involved in gluing the two parts together. 3 Application Processors Checkpointing Themselves Here the goal is for an application processor to checkpoint its state in such a way that if a rollback is called for, due to the failure of another processor, the processor can roll back to its most recent checkpoint. In standard disk-based systems, a processor checkpoints itself by saving the contents of its address space to disk. Typically this involves saving all values in the stack, heap, global variables and registers as in Figure 1(a). If the processor must roll back, it overwrites the current contents of its address space with the stored checkpoint. As a last step, it restores the registers, which restarts the computation from the checkpoint, thereby completing the rollback. For more detail on general process checkpointing and recovery, see the papers on Condor [34] and libckpt [25]. Memory Disk Registers Application Processor Address space (unused) Memory Registers Application Processor Address space Diskless checkpoint Memory Registers Application Processor Address space Page faults Memory Clone Registers Application Processor Address space Checkpoint (unused) (a) (b) (c) (d) Figure 1: (a) Checkpointing to disk, (b) simple diskless checkpointing, (c) incremental diskless checkpointing, (d) forked diskless checkpointing With diskless checkpointing, the processor saves its state in memory, rather than on disk. In its simplest form, diskless checkpointing requires an in-memory copy of the address space and registers, as in Figure 1(b). If a rollback is required, the contents of the address space and registers are restored from the in-memory checkpoint. Note that this checkpoint will not tolerate the failure of the application processor itself; it simply enables the processor to roll back to the most recent checkpoint if another processor fails. One drawback of simple diskless checkpointing is memory usage. A complete copy of the application must be retained in the memory of each application processor. A solution to this problem is to use incremental checkpointing [11, 38], as in Figure 1(c). To take a checkpoint, an application processor sets the virtual memory protection bits of all pages in its address space to be read-only [1]. When the application attempts to write a page, an access violation (page fault) occurs. The checkpointing system then makes a copy of the faulting page, and resets the page's protection to read-write. Thus, a processor's checkpoint consists of the read-only pages in its address space plus the stored copies of all the read-write pages. To roll back to a checkpoint, the processor simply copies (or maps) the checkpointed copies of all its read-write pages back to the application's address space. As long as the application does not overwrite all of its pages between checkpoints, incremental checkpointing improves both the performance and memory utilization of checkpointing. The last useful checkpointing method is forked (or copy-on-write) checkpointing [9, 21, 25]. To checkpoint, the application clones itself (with, for example, the fork() system call in Unix) as depicted in Figure 1(d). This clone is the diskless checkpoint. To roll back, the application overwrites its state with the clone's, or if possible, the clone merely assumes the role of the application. Forked checkpointing is very similar to incremental checkpointing because most operating systems implement process cloning with copy-on-write. This means that the process and its clone will share pages until one of the processes alters the page. Thus, it works in the same manner as incremental checkpointing, except the identification of modified pages and the page copying are all performed in the operating system. This results in less CPU activity switching back and forth from system to user mode. Moreover, forked checkpointing does not require that the user have access to virtual memory protection facilities, which are not available in all operating systems. 4 Encoding the checkpoints The goal of this part is for extra checkpoint processors to store enough information that the checkpoints of failed processors may be reconstructed. Specifically, there are m checkpoint processors. These processors encode the checkpoints of the application processors in such a way that when application processors fail, their checkpoints may be recalculated from the checkpoints of the non-failed processors plus the encodings in the checkpoint processors. 4.1 Parity (Raid Level 5) The simplest checkpoint encoding is parity (Figure 2(a)). Here there is one checkpoint processor (i.e. that encodes the bitwise parity of each of the application's checkpoints. In other words, let byte b j represent the j-th byte of application processor i. Then the j-th byte of the checkpoint processor will be: If any application processor fails, the state of the system may be recovered as follows. First, a replacement processor is selected to take the place of the failed application processor. This could be the checkpoint processor, a spare processor that had previously been unused, or the failed processor itself if the failure was transient. The replacement processor calculates the checkpoint of the failed processor by taking the parity of the checkpoints of the non-failed processors and the encoding in the checkpoint processor. In other words, suppose processor i is the failed processor. Then its checkpoint may be reconstructed as: ckp Note that this is the same recovery scheme as Raid Level 5 in disk array technology [5]. When the replacement processor has calculated the checkpoint of the failed processor, then all application processors roll back to the previous checkpoint, and the computation proceeds from that point. Besides parity, there are several other schemes than can be used to encode the checkpoints. They vary in the number of checkpoint processors, the efficiency of encoding, and the amount of failure coverage. They are detailed below. (a) (b) (c) Figure 2: Encoding the checkpoints: (a) Raid Level 5, (b) Mirroring, (c) One-dimensional parity 4.2 Mirroring Checkpoint mirroring (Figure 2(b)) is another simple encoding scheme. With mirroring, there are checkpoint processors, and the i-th checkpoint processor simply stores the checkpoint of the i-th application processor. Thus, up to n processor failures may be tolerated, although the failure of both an application processor and its checkpoint processor cannot be tolerated. Checkpoint mirroring should have a very low checkpointing overhead because no encoding calculations (such as parity) need to be made. 4.3 1-dimensional parity With one-dimensional parity (Figure 2(c)) there are checkpoint processors. The application processors are partitioned into m groups of roughly equal size. Checkpoint processor i then calculates the parity of the checkpoints in group i. This increases the failure coverage, because now one processor failure per group may be tolerated. Moreover, the calculation of the checkpoint encoding should be more efficient because there is no longer a single bottleneck (the checkpoint processor). Note that 1-dimensional parity reduces to Raid Level 5 when and to mirroring when 4.4 2-dimensional parity Two-dimensional parity (Figure 3(d)) is an extension of one-dimensional parity. With two-dimensional parity, the application processors are arranged logically in a two-dimensional grid, and there is a checkpoint processor for each row and column of the grid. Each checkpoint processor calculates the parity of the application processors in its row or column. Two-dimensional parity requires m - 2 checkpoint processors, and can tolerate the failure of any one processor in each row and column. This means that any two-processor failures may be tolerated. (d) (e)2 a Figure 3: Encoding the checkpoints: (d) Two-dimensional parity, (e) Hamming coding, (f) EvenOdd coding, (g) Reed-Solomon coding 4.5 Other parity-based codes The well-known Hamming codes (Figure 3(e)) may be used to tolerate any two-processor failures with the addition of roughly log n processors [13]. Each checkpoint processor calculates the parity of a subset of the application processors. EvenOdd coding (Figure 3(f)) is a technique where checkpoint processors are employed and all two-processor failures may be tolerated [3]. The encoding is based on parity calculations, but is a little more complex than the above schemes. 4.6 Reed-Solomon coding The most general purpose encoding technique is Reed-Solomon coding [24] (Figure 3(g)). Here m checkpointing processors use Galois Field arithmetic to encode the checkpoints in such a way that any m failures may be tolerated. Since the encoding is more complex than parity, the CPU overhead of Reed-Solomon coding is greater than the other methods, but it achieves maximal failure coverage per checkpoint processor. 5 Gluing the two parts together Sections 3 and 4 have discussed how application processors store checkpoints internally, and how the checkpoint processors encode information. The final component of diskless checkpointing is coordinating the application and checkpointing processors in an efficient and correct way. This section discusses the relevant details in the coordination of the two sets of processors. We focus primarily on Raid Level 5 encodings, and then discuss the differences that the other encodings entail. 5.1 Tolerating failures when checkpointing As with all checkpointing systems, diskless checkpointing systems must take care to remain fault-tolerant even if there is a failure while checkpointing or recovery is underway. This is done by making sure that each coordinated checkpoint remains valid until the next coordinated checkpoint has been completed. The checkpointing processors control this process. When all the checkpointing processors have completed calculating their encodings for the current checkpoint, then they may discard their previous encodings, and then notify the application processors that they may discard their previous checkpoints. Upon recovery, if the checkpointing processors all have valid encodings for the most recent checkpoint, then these are used for recovery, along with the most recent checkpoints in the non-failed application processors. If any checkpointing processor does not have a valid encoding for the most recent checkpoint, then the previous encoding must be used along with the previous checkpoints in the non-failed application processors. This protocol ensures that there is always a valid coordinated checkpoint of the system in memory. If all checkpoint processors have their encodings for coordinated checkpoint i, then all application processors will have their checkpoints for coordinated checkpoint i. If any checkpoint processor has an incomplete encoding for checkpoint i, then all checkpoint processors will still contain their encodings for coordinated checkpoint Moreover, all application processors will have their checkpoints for coordinated checkpoint i \Gamma 1. Thus, the whole system may recover to coordinated checkpoint If a failure is detected during recovery, then the remaining processors simply initiate the recovery procedure anew. 5.2 Space demands A ramification of the preceding protocol is that at the moment when the checkpoint processors finish storing their encodings, all processors contain two checkpoints in memory: the current checkpoint and the previous checkpoint. Thus, the memory usage of diskless checkpointing is a serious issue. Suppose the size of an application processor's address space is M bytes. Then simple diskless checkpointing consumes an extra M bytes of memory to hold a checkpoint. To ensure that only M bytes of extra memory are consumed at all times, the application must be frozen during checkpointing. Then the application's address (a) (b) Figure 4: Calculating the encoding: (a) direct, (b) fan-in space may be used (without being copied) to calculate the checkpoint encodings. When the encodings have been calculated, the application's address space may be copied over its previous checkpoint, which is now expendable. Then the application is unfrozen. With incremental checkpointing, checkpointed copies of pages are made when page faults are caught. At checkpoint time, the processors calculate the encodings, then discard the checkpointed copies of pages and set the protection of all application pages to read-only. Thus, if the incremental checkpoint size is I, then only I extra bytes of memory are necessary. In the worst case, all pages are modified between checkpoints, and I equals M . With forked checkpointing, each checkpoint is a separate process. When the checkpoint processors complete their encodings, there are three processes contained by each application processor: the application itself, its most recent checkpoint, and the previous checkpoint. Since process cloning uses the copy-on-write optimization, each checkpoint process only consumes an extra I bytes of memory. Therefore, forked checkpointing requires an extra 2I bytes of memory during checkpointing, and I bytes at all other times. In the worst case, this is 2M during checkpointing, and M at other times. Finally, disk-based checkpointing using the fork optimization requires I 0 bytes of memory, where I 0 consists of all pages that are modified while checkpointing is taking place. I 0 should be less than I, though if the latency of checkpointing is large compared to the checkpointing interval, I 0 may be close to I. 5.3 Sending and calculating the encoding With Raid Level 5 encoding, there is one checkpoint processor C 1 , and n application processors stores the bitwise parity of the checkpoints of each application processor. The simplest way to calculate the parity is to employ the direct method: each application processor simply sends its checkpoint to C 1 . Initially, clears a portion of its memory, which we call e 1 , to store the checkpoint encoding. Upon receiving ckp i from This is shown in Figure 4(a). In Figure 4, the \Phi signs are shown directly above the processors that perform the bitwise exclusive or. Arrows from one processor to another represent one processor sending its checkpoint to another. There are two problems with the direct method. First, C 1 can become a message-receiving bottleneck, since it is the destination of all checkpoint messages. Second, C 1 does all of the parity calculations. Both problems may be alleviated with the fan-in algorithm. Here, the application processors perform the parity calculation in log n steps and send the final result to C 1 , which stores the result in its memory. This is shown in Figure 4(b). For other encodings besides Raid Level 5, these two methods may be extended. In the direct method, each processor sends its checkpoint in a multicast message to the proper checkpointing processors. If necessary (e.g., for Reed-Solomon coding), the checkpointing processors modify the checkpoints, and then exclusive-or them into their checkpoints. In the fan-in method, there is one fan-in performed for each checkpointing processor. This may entail the cooperation of all application processors (e.g., in Reed-Solomon coding), or a subset of the application processors (e.g., in one-dimensional parity). If a checkpoint must be modified for the encoding, it is done at application processor P i before the fan-in starts. For most networks, the fan-in algorithm will be preferable to the direct because it eliminates bottlenecks and distributes the parity calculations. However, if the network supports multicast, the encodings involving multiple checkpointing processors may profit from the direct method. 5.4 Breaking the checkpoint into chunks The preceding description implies that whole checkpoints are sent from processor to processor. Since checkpoints may be large, it often makes more efficient use of memory to break the checkpoint into chunks of a fixed size. For example, in the fan-in algorithm, only two extra chunks of memory are needed to receive an incoming chunk from another processor, make the parity calculation, and then send off the result. The chunks should be small enough that they do not consume too much memory, but large enough that the overhead in sending chunks is not dominated by message-sending start-up. 5.5 Sending diffs If the application processors use incremental checkpointing, then they can avoid overhead by sending only pages that have been modified since the previous checkpoint. However, this can cause problems in creating the checkpoint encoding. Specifically, if the encoding is to be created anew at each checkpoint, it needs to have all checkpointing data from all processors. The solution to this is to use diffs. Assume that the direct encoding method is being employed. The checkpoint processor first copies its previous checkpoint to its current checkpoint. Then each application processor does the following. For each modified page page k in its address space, it calculates diff k , which is the bitwise exclusive-or of the current copy of the page and the copy of the page in the previous checkpoint (which of course is available to the application processor). It then sends diff k to the checkpoint processor, which XOR's it into its checkpoint. This has the effect of subtracting out the old copy of the page and adding in the new copy. In this way, unmodified pages need not be sent to the checkpointing processor. One may use diffs with the fan-in algorithm as well, stipulating that if a processor does not modify a page during the checkpoint interval, then it does not need to send that page or XOR it with other pages when performing the fan-in. 5.6 Compressing Diffs By sending diffs rather than actual bytes of the checkpoint, an interesting opportunity for compression arises. Suppose that an application modifies just a few bytes on a page. Then the diff of that page and its previously checkpointed copy will be composed of mostly zeros, which can be easily compressed using either run-length encoding or an algorithm that sends tagged bytes rather than whole pages. Such compression trades off use of more CPU for a reduced load on the network. Compression combines naturally with incremental checkpointing, where modified pages are compressed before being sent. It may also be used with simple and forked checkpointing by converting the entire checkpoint into a diff and compressing it before sending it along. This has the effect of emulating incremental checkpointing, because regions of memory that have not been modified get compressed to nothing. 6 Implementation and Experiment In order to assess the performance of diskless checkpointing as compared to standard disk-based checkpointing on networks of workstations, we implemented a small transparent checkpointing system on a network of 24 Sun Sparc5 workstations at the University of Tennessee. Each workstation has 96 Mbytes of physical memory and runs SunOS version 4.1.3. The workstations are connected to each other by a fast, switched Ethernet which can be isolated for performance testing. The measured peak bandwidth between any two processors is roughly 5 megabytes per second. The workstations have very little accessible local disk storage: 38 megabytes per machine. However, the machines are connected via regular Ethernet to the department's file servers using Sun NFS. These disks have a bandwidth of 1.7 megabytes per second, but the performance of NFS on the Ethernet is far worse. With NFS, remote file writes achieve a bandwidth of 0.13 megabytes per second. The page size of each machine is 4096 bytes, and access to the page tables is controlled by the mprotect() system call. Our checkpointer runs on top of PVM [12] and works like many PVM checkpointers [4, 33]. Applications do not need to be recompiled, but their object modules must be relinked with our checkpointing/modified PVM library. When the applications are started, the checkpointing code gets control and reads startup information from a control file. This information includes the checkpoint interval, which checkpointing optimizations to use, plus where checkpoints should be stored (to disk or to checkpointing processors). The application then starts, and one of the application processors is interrupted when the checkpointing interval has expired. This processor coordinates with the other application processors using the "Sync-and-stop" synchronization algorithm, and once consistency has been determined, the processors checkpoint. Abbreviation Description checkpointing DISK-FORK Checkpointing to disk using fork() SIMP Simple diskless checkpointing INC Incremental diskless checkpointing Forked diskless checkpointing INC-FORK Incremental, forked diskless checkpointing C-SIMP Simple diskless checkpointing with compression C-INC Incremental diskless checkpointing with compression Forked diskless checkpointing with compression C-INC-FORK Incremental, forked diskless checkpointing with compression Table 1: Checkpointing variants implemented in our experiments PVM includes some basic forms of failure detection. Specifically, if a processor in the current PVM session fails, the rest of the processors eventually notice the failure and remove the failed processor from the PVM session. PVM allows the user to be notified of such events. Our checkpointer uses this facility to recognize processor failures. When such a failure occurs, then if there is a spare processor in the PVM session, it is selected to replace the failed processor. If there is no spare processor, and diskless checkpointing is being employed, then a checkpoint processor is chosen to be the replacement processor. Recovery proceeds automatically, either from the disk-based or diskless checkpoint. It is important to note that our checkpointer does not require the programmer to modify his or her code to enable checkpointing. A simple relinking is all that is necessary. The gamut of checkpointing variants is enumerated in Table 1. This includes standard disk-based checkpointing using the fork() optimization. We do not test incremental, disk-based checkpointing because it does not improve the performance of checkpointing in any of our tests. 1 . For diskless checkpointing, we implement Raid Level 5 encoding using the fan-in algorithm. Checkpoint encodings are created in chunks of 4096 bytes (conveniently, also the page size). The choice of algorithm has some ramifications on how certain optimizations work. For example, when performing incremental checkpointing, the encoding is created chunk-by-chunk, but if a processor has not modified the corresponding page, then an empty message is sent as part of the fan-in instead of the page. When using diff-based compression, pages are compressed using a bitmap-based compression algorithm [29]. Compression is performed by the sending processor before sending, and then uncompressed by the receiving 1 This is not to say that incremental, disk-based checkpointing is not often a useful optimizations. It simply does not help in any of our tests. Application Running Time Checkpoint Size per node (sec) (h:mm:ss) (Mbytes) NBODY 5722 1:35:22 3.7 CELL 6351 1:45:51 41.4 PCG 5873 1:37:53 66.6 Table 2: Basic parameters of the testing applications processor, which merges the page with its own, and compresses the result before sending it along. When the final compressed chunk reaches the checkpointing processor, it uncompresses the chunk and merges it with the previous checkpoint encoding, which is then stored as the next encoding. Applications We used five applications to test the performance of checkpointing. These applications are all CPU-intensive, parallel programs of the sort that often require hours, or sometimes days of execution. We executed instances of these programs that took between 1.5 and 2 hours to run on sixteen processors in the absence of checkpointing. In all cases, it is clear how the programs scale in size, and how this scaling will affect the performance of checkpointing. The basic parameters of each application are presented in Table 2. We briefly describe each application, ordered by checkpoint size, below. 7.1 NBODY NBODY computes N-body interactions among particles in a system. The program is written in C, and uses the parallel multipole tree algorithm [19]. The instance used in our tests was 15,000 particles and ten iterations. The basic structure of the program is as follows. Each particle is represented by a data structure with several fields. The particles are partitioned among "slave" processors (sixteen in our tests) in such a way that processors that are "close to each other" (by some metric) reside in the same slave, to limit interslave communication. For this reason, slave processors can differ in the number of particles they hold and therefore in their sizes. For example, in our tests, the slave processors averaged 3.7 megabytes in size, but the largest was six megabytes. At each iteration, the "location" field (among others) of each particle is updated to reflect the n-body inter- action. Since the size of a particle's data structure is less than the machine's page size, this means that almost all pages of the slave processors are modified during each iteration, leading to poor incremental checkpointing behavior when the checkpointing interval spans multiple iterations. However, since much of each particle's data is left unmodified from iteration to iteration, only a few bytes per page are changed, resulting in good diff-based compression. There are two parameters that affect the running time and memory usage of NBODY. These are the number of particles, which affects both time and space, and the number of iterations, which only affects the running time. NBODY is the only application where the checkpoints are small enough to allow the same number of checkpoints in both diskless and disk-based checkpointing. 7.2 MAT MAT is a C program that computes the floating point matrix product of two square matrices using Cannon's algorithm [17]. The matrix size in our tests was 4,608\Theta4,608, leading to 15.1 megabyte checkpoints per processor. On a uniprocessor, matrix multiplication typically shows excellent incremental checkpointing behavior, since the two input matrices are read-only, and the product matrix is calculated sequentially, filling up whole pages at a time in such a way that once a product element is calculated, it is never subsequently modified [25]. However, most high-performance parallel algorithms, such as Cannon's algorithm, differ in this respect. In Cannon's algorithm, all three matrices are partitioned in square blocks among the n processors (and it is assumed n is a perfect square). The algorithm proceeds in p steps. In each step, each processor adds the product of its two input submatrices to its product submatrix. Then the processors send their input submatrices to neighboring processors, receiving new ones in their place, and repeat until the product submatrices are calculated. The ramification of this data movement is that during the course of an iteration, all matrices are modified. Therefore, if checkpoints span iterations (as is the case in disk-based checkpointing), incremental checkpointing will have no beneficial effect. If multiple checkpoints are taken in the same iteration (as is the case in diskless checkpointing), then incremental checkpointing will be successful as in the uniprocessor case. When pages are updated in MAT, they are updated in their entirety, leading to very poor diff-based compression MAT's time and space demands are determined by the size of the matrix. For an N \Theta N matrix, the memory usage is proportional to N 2 , and the running time is proportional to N 3 . The communication patterns of MAT depend on the number of processors, and are the same for all matrix sizes. MAT and NBODY are the only applications where it is possible to take more than one disk-based checkpoint during the program's execution. Three disk-based checkpoints (as opposed to seven diskless checkpoints) are taken in MAT. 7.3 PSTSWM PSTSWM is a fortran program that solves the nonlinear shallow water equations on a rotating sphere using the spectral transform method [14]. The instance used here simulates the state of a 3-D system for a duration of hours. Like NBODY, PSTSWM modifies the majority of its pages during each iteration, but it only modifies a few bytes per page. Therefore, incremental checkpointing should show limited improvement, but diff-based compression should work well. PSTSWM's checkpoints are large - approximately 25 megabytes per processor. However, since each machine has 96 megabytes of physical memory, two checkpoints may be stored in their stressing the limits of physical memory. PSTSWM can scale in size by simulating a denser particle grid. Once the size is set, each iteration performs roughly the same actions. Therefore, simulating longer time frames increases the running time in a linear fashion without altering the general behavior (e.g. memory access pattern) significantly. 7.4 CELL CELL is a parallel cellular automaton simulation program. Written in C, this program distributes two grids of cellular automata evenly across all the application processors. One grid is denoted current, and one is denoted next. The values of the current grid are used to calculate the values in the next grid, and then the two grids' identities are swapped. The instance used in our tests simulates a 18,512 by 18,512 cellular automaton grid for generations. During each iteration, CELL updates every automaton in the next grid. Therefore, if checkpoints span two or more iterations, all memory locations will be updated, rendering incremental checkpointing useless. Compressibility depends on the data itself. "Sparse" grids (where many automata take on zero values) may see little change in the automata's values over time, which can lead to good compression. Denser grids lead to less compression. In our tests, we used very sparse grids. The program size is directly proportional to the grid size. The running time is proportional to the grid size times the number of iterations. Each pair of iterations performs the same operations, and thus has the same memory access and communication patterns. PCG is a fortran program that solves for a large, sparse matrix A using the "Preconditioned Conjugate Gradient" iterative method The matrix A is converted to a small, dense format, and then approximations to x are calculated and refined iteratively until they reach a user-specified tolerance from the correct values. In our tests, A is a 1,638,400 by 1,638,400 element sparse matrix, and the program takes 3,750 iterations. The exact mechanics and memory usage of PCG are detailed in [26]. The salient points are as follows. The main data structures in the program may be viewed as many vectors of length N (in our instances, These vectors are distributed among all the application processors. Roughly three quarters of these vectors are never modified once the program starts calculating. The rest are updated in their entirety at each iteration. Therefore, incremental checkpoints should be one quarter the size of non-incremental checkpoints. The data that gets updated at every iteration is stored densely on contiguous pages, offering little opportunity for diff-based compression. The program size is directly proportional to N , and like CELL and PSTSWM, the running time is proportional to the size times the number of iterations. Each application processor holds 66.6 megabytes worth of data in PCG. Therefore, one simple diskless check-point will not fit into memory. However, when incremental and copy-on-write checkpointing are employed, the application and one or two checkpoints consume just a few megabytes more memory than is available. The size of the checkpoints combined with the speed of Sun NFS results in the inability to take disk-based checkpoints of PCG. This is because the time to store one checkpoint is longer than the running time of the application. It should be reiterated that the instances for these tests were chosen to run for a period of time that was long enough to measure the impact of checkpointing and recovery. In all applications, there are natural input parameters which result in longer execution times and larger checkpoints. Our goal in these tests is to assess the performance of checkpointing so that users of longer-running applications may be able to project the expected running time of their applications in the presence of failures while employing the various checkpointing variants. The raw data for the experiments is in the Appendix of this paper. All graphs in this section are derived directly from the raw data. In most cases, the tests were executed in triplicate. The number of times each test was executed plus the standard deviations in execution times is displayed in the tables in the Appendix. The tables and graphs display average data. We concentrate on two performance measures: latency and overhead. Latency is the time between when a checkpoint is initiated, and when it may be used for recovery. Overhead has been defined previously. Overhead is a direct measure of the performance penalty induced on an application due to checkpointing. The impact of latency is more subtle, and will be discussed in detail in Section 9. 8.1 Checkpointing to disk Figure 5 plots checkpoint latency and overhead of checkpointing to disk (the DISK-FORK tests). These are plotted as a function of the applications' per-processor checkpoint sizes. As displayed in leftmost graph, the latency in the DISK-FORK tests is directly proportional to the checkpoint size, achieving a bandwidth of Mbytes/sec. Here bandwidth is calculated as per-processor checkpoint size times the number of processors, divided by the checkpoint latency. Using that information, the checkpoint latency of the PCG test is projected to be roughly 8,663 seconds. The rightmost graph displays overhead as a function of checkpoint size. While the graph appears roughly linear, it should be noted that the overhead of checkpointing is not a simple function of checkpoint size. The bulk Checkpoint size (Mbyte / processor)30009000 Latency per Checkpoint (sec) MAT PCG (projected) Checkpoint size (Mbyte / processor)3090 per Checkpoint (sec) MAT Figure 5: Checkpoint latency and overhead of checkpointing to disk (DISK-FORK) Checkpoint size (Mbyte / processor)100300Latency Per Checkpoint (sec) Checkpoint size (Mbyte / processor)100300Overhead per Checkpoint (sec) Figure Checkpoint latency and overhead of SIMP and FORK of work performed in checkpointing involves DMA from each processor's memory to its network interface card. The CPU is only affected significantly when one of the following occurs: ffl A DMA transaction needs to be initiated or repeated. ffl A copy-on-write page fault occurs in the application. ffl There is contention for the memory bus. There are also effects on the cache as a result of checkpointing. Therefore, although checkpoint size is a rough measuring stick for computing the overhead of DISK-FORK checkpointing, it is not the whole story. As has been shown in other research, the copy-on-write optimization does an excellent job of reducing overhead [9, 21, 25]. In this test, the overhead is between 0.7 and 5.5 percent of the checkpoint latency. 8.2 Diskless checkpointing: SIMP and FORK Figure 6 plots checkpoint latency and overhead of the SIMP and FORK tests, again plotted as a function of checkpoint size. As in the DISK-FORK case, both the SIMP and FORK latencies are directly proportional to the checkpoint size, with the exception of the SIMP test in the PCG application. Here, the combined size of the application and its checkpoint exceeds the size of physical memory, resulting in pages being swapped to the backing store. This degrades the performance of checkpointing. In the FORK test, the checkpoint only requires an additional 16.6 Mbytes of memory, since the unmodified pages of memory are shared between the application and its checkpoint. Therefore, the checkpoint latency follows the same linear pattern as in the other applications. With the exception of the SIMP test in the PCG application, the bandwidth of checkpointing in SIMP and FORK is roughly 4.4 Mbytes/sec. This is a factor of 34 faster than the DISK-FORK bandwidth. The overhead of the SIMP tests is identical to the latency, since the application is halted during checkpointing. In the FORK tests, the overhead is reduced by 29.4 (in MAT) to 53.7 (in PCG) percent. Although this is an improvement, it is not the same degree of improvement as in the DISK-FORK tests. The reason for this is that the CPU is more involved in diskless checkpointing than in disk-based checkpointing. In diskless checkpointing, the parity of each processor's checkpoint must be calculated, and this takes the CPU (plus some memory) away from the application. The only time when disk-based checkpointing makes more use of the CPU than diskless checkpointing is when the longer latency of checkpointing causes more copy-on-write page faults to occur. 8.3 The rest of the tests All of the diskless checkpointing results are displayed in Figure 7. The top row of graphs shows the checkpoint latency for each test in each application. The middle row shows checkpoint overhead, and the bottom row shows the average checkpoint size. This is a bit of a misnomer, because in all cases, the in-memory and parity processor checkpoints are the same size. However, with incremental checkpointing and compression, fewer bytes are sent per processor. The "checkpoint size" graphs (and the "checkpoint size" columns in the Appendix) display the average number of bytes that each processor sends during checkpointing. Some salient features from Figure 7 are as follows. First, incremental checkpointing significantly reduces the average checkpoint sizes in the MAT and PCG applications. In the other three applications, the checkpoint size of SIMP and INC are roughly the same. In the MAT and PCG applications, significant reductions in checkpoint latency and overhead result from incremental checkpointing. In both cases, the mixture of incremental and forked checkpointing result in the lowest overhead of the all diskless checkpointing tests. When incremental checkpointing fails to decrease the size of checkpoints, as in the NBODY and CELL ap- plications, the overhead of checkpointing is greater than with simple checkpointing. In both of these applications, the INC-FORK tests yielded the highest checkpoint latencies. The results of diff-based compression are interesting. In three applications (NBODY, PSTSWM and CELL), Checkpoint Latency (sec) MAT INC-FORK C-SIMP C-FORK C-INC C-INC-FORK515Checkpoint Overhead (sec) MAT Checkpoint Size MAT Figure 7: Diskless checkpoint latency, overhead, and size per application incremental checkpointing fails because most of the programs' pages are updated at every iteration. However, diff- based compression succeeds in reducing checkpoint size because the pages are either sparsely modified (NBODY and PSTSWM) or updated with the same values (CELL). In these three applications, the C-FORK tests yielded the lowest checkpoint overhead. Note that since compression adds extra demands on the CPU, the reduction in overhead is not as drastic as with incremental checkpointing. It is also interesting to note that the lowest overhead is achieved with C-FORK rather than C-INC or C-INC-FORK. This is because in these tests, almost all pages are modified between checkpoints, and therefore incremental checkpointing merely adds the overhead of processing page faults. In the other two tests (MAT and PCG), diff-based compression brings the checkpoint sizes of the FORK and SIMP tests to roughly the same size as incremental checkpointing. However, it does not improve upon incremental Application Recovery Time (sec) PSTSWM 66.3 CELL 138.3 PCG 375.3 Table 3: Recovery times for the SIMP tests checkpointing in terms of size or overhead. This is because the modified pages showed little compressibility. 8.4 Recovery time Table 3 shows the time that it takes the system to recover from a single failure and continue execution from the most recent checkpoint during the SIMP tests. Here, a processor failure is simulated by terminating one of the application processors. PVM has been written so that the other processors recognize this failure, and our modifications take advantage of this to automate the process of recovery. In our tests, the checkpointing processor takes the place of the failed application processor. The recovery times are roughly equal to the checkpoint latencies of the SIMP applications. It should be noted that in all but the DISK-FORK tests, the recovery times are equal, since the entire diskless checkpoint of the failed processor must be calculated. In the DISK-FORK tests, the recovery times are equal to the checkpoint latencies. Thus, like the latencies, they are extremely large. 9 Discussion 9.1 Diskless vs. disk-based checkpointing There are two basic results that we may draw from our tests concerning diskless vs. disk-based checkpointing: ffl The checkpoint latency and recovery time of diskless checkpointing is vastly lower than disk-based checkpointing. As stated in section 8.2, the latency (and recovery time) of disk-based checkpointing is a factor of 34 slower than diskless checkpointing. This is a result of the poor performance of Sun NFS combined with the fact that all processors use the same disk. ffl The overhead of diskless checkpointing is comparable to disk-based checkpointing. Figure 8 plots the overhead of disk-based checkpointing and the overhead of the best diskless variant for each application. Checkpoint Overhead (sec) Figure 8: Checkpoint overhead of disk-based checkpointing as compared to the best diskless variant. In some cases (NBODY and PSTSWM), diskless checkpointing outperforms disk-based, and in others disk-based outperforms diskless. The question mark is plotted in PCG because we were unable to complete a disk-based checkpoint during the lifetime of the application. There are two reasons why diskless checkpointing may be viewed as preferable to disk-based checkpointing. First, it lowers the expected running time of the application in the presence of failures. Second, it has less effect on the computing environment, which is of special concern if the environment is shared. We consider each of these in turn. 9.1.1 Expected running time Supposing that failure rate is governed by a Poisson process, Vaidya has derived equations for assessing the performance of an application in the presence of checkpointing and rollback recovery [36]. These equations take as input the average overhead, latency, and recovery time per checkpoint, plus the rate of failures, and are defined as follows. where: The rate of failures (1=MTBF ). The optimal checkpoint interval. The average overhead per checkpoint. The average latency per checkpoint. recovery time from a checkpoint. The running time of the application in the absence of checkpointing, recovery, and failures (i.e. the BASE test). The "overhead ratio," which is a measure of the performance penalty due to checkpointing, recovery and failures[36]. The expected running time of the optimal checkpoint interval in the presence of failures, checkpointing and recovery. The optimal expected running time of the application in the presence of failures, checkpointing and recovery. The expected running time of the application in the presence of failures, but no checkpointing and recovery (i.e. the application is restarted from scratch following a failure). In all these equations, the repair time is assumed to be zero. This approximates the case when a spare processor is ready to continue computation immediately following a failure. If repair time is significant, then Eq's 2 and 5 become: These equations may be used to compare checkpointing algorithms as follows. First, for each algorithm T opt may be calculated from - and O using Eq. 1. Next, \Gamma and r may be determined by Eqs. 2 and 3. If so desired, the expected running time of an application (T ckp ) for each algorithm may then be determined by Eq. 4. The checkpointing algorithm with the lowest value of r will be the one with the smallest expected running time. Thus, r suffices as a metric by which to compare checkpointing algorithms. If T ckp is greater than T nockp , then the application cannot benefit from checkpointing. This occurs when the application's running time (T base ) is not significantly greater than T opt . However, as T base grows, T nockp increases more rapidly than T ckp to the point that checkpointing improves the program's expected running time in the presence of failures. In Table 4, we use the data from Section 8 to derive values for T opt , \Gamma, r, T ckp and T nockp for each of the tests presented in Figure 8. We calculated - in the following manner. In their study of host reliability on the Internet, Long et at [22] determined an average MTBF of 29.29 days. Assuming independent processor failures, this means that the MTBF of a collection of 16 processors is days, and the MTBF of a collection of 17 processors is days. This gives - a value of 6:301 10 \Gamma6 failures per second for and 6:694 10 \Gamma6 failures per second for 17 processors. We use the former value as the failure rate for disk-based checkpointing and for no checkpointing, and the latter value for diskless checkpointing. Table 4 shows that in all applications, diskless checkpointing performs better than disk-based checkpointing. This can be seen in the lower expected running times (T ckp ), and the lower overhead ratios (r). Therefore, even though the two have similar checkpoint overheads, the extremely large latency and recovery time of disk-based checkpointing makes it unattractive in comparison to diskless checkpointing. Another significant result of Table 4 is that in two applications, NBODY and MAT, the expected running time in the presence of failures is minimized by diskless checkpointing. In the other three applications, no checkpointing Application Test Tbase Topt \Gamma r Tckp Tnockp (sec) (sec) (sec) (sec) (sec) Table 4: Calculated values of T int , \Gamma, r, T ckp and T nockp . gives the smallest expected running time. That any checkpointing improves performance is somewhat surprising, given the relatively small execution times of the experiments with respect to the MTBF. There are no cases where disk-based checkpointing gives a smaller expected running time. As the execution time of an application grows, checkpointing becomes much more attractive. For example, suppose the user desires to simulate 5000 hours in PSTSWM instead of 102. Then the program will take roughly 275,000 seconds, or 3.18 days. Such an execution would not alter the size of the checkpoints, and therefore we may use the same overhead, latency and recovery times as presented in Section 8. This leads to expected execution times of 3.256 days for diskless checkpointing, 3.390 days for disk-based checkpointing and 8.553 days for no checkpointing. 9.1.2 The effect on shared resources Large checkpoint latencies can be detrimental in other ways. For example, in disk-based checkpointing, the entire latency period is spent writing checkpoint data to stable storage. If other programs or users share the stable storage, large checkpoint latencies are undesirable, because the performance of stable storage as seen by others is degraded for a long period of time. In [23], the effect of DISK-FORK checkpointing on the performance of stable storage was assessed. While a checkpoint was being stored to the central disk, a processor not involved in the application timed the bandwidth of disk writes. In that test, the performance of stable storage was degraded by 87 percent. This is significant, for it means that extremely long checkpoint latencies, such as those measured in our tests, have the potential to degrade the performance of the system in a severe manner for a long time. Diskless checkpointing, on the other hand, exhibits much smaller checkpoint latencies, and because the calculation of the checkpoint encoding involves both the network and the CPU, the impact on shared resources (in this case, the network) is far less [23]. 9.2 Recommendations Given the results of these experiments, we can make the following recommendations. Of the checkpointing variants tested in this paper, three stand out as the most useful: DISK-FORK, C-FORK and INC-FORK. On a system with similar performance to ours, each is the most useful in certain cases: ffl If checkpoints are small or the likelihood of wholesale system failures is high, then DISK-FORK checkpointing should be employed. ffl If the program modifies a few bytes per page between checkpoints, or if the machine does not provide access to virtual memory facilities, then C-FORK diskless checkpointing should be employed. ffl If the program does not modify a significant number of pages between checkpoints, then INC-FORK diskless checkpointing should be employed. Although we did not test such applications, there may be times when FORK and SIMP are the most useful checkpointing methods. This is when all pages are modified in a dense manner between checkpoints. Then FORK will have the lowest overhead when there is enough memory to store two checkpoints, and SIMP will have the lower overhead otherwise. None of our applications would have benefited from incremental checkpointing to disk. However, if multiple checkpoints are taken and the program modifies only a fraction of its pages between checkpoints, incremental forked checkpoints will outperform DISK-FORK. Finally, in interpreting the results, it is important to note that the speed of stable storage in these experiments is quite slow. A faster network, a faster file system, or a file system with multiple disks will improve the performance of disk-based checkpointing relative to diskless checkpointing. On the other hand, a system with more processors will degrade the performance of disk-based checkpointing relative to diskless checkpointing. It should be possible using the equations in Section 9.1.1 to extrapolate the results of our experiments to systems with different performance parameters. Related Work There has been much research performed on checkpointing and rollback recovery. The important algorithms and performance optimizations for disk-based checkpointing in parallel and distributed systems are presented in [8]. Research more directly related to diskless checkpointing is cited below. The first paper on diskless checkpointing was presented by Plank and Li [27]. This paper may be viewed as a completion of that original paper. Silva et al [32] implemented checkpoint mirroring on a transputer network, and performed experiments to determine that it outperformed disk-based checkpointing. Chiueh and Deng [6] implemented checkpoint mirroring and Raid Level 5 checkpointing on a massively parallel (4096 processors) SIMD machine. They found that mirroring improved performance by a factor of 10. Both implementations involved modifying the application to perform checkpointing, rather than simply relinking with a checkpointing library. Scales and Lam [31] implemented a distributed programming system built on special primitives with shared-memory semantics. They use redundancy built into the system, plus checkpoint mirroring when necessary to tolerate single processor failures with low overhead. In a similar manner, Costa et al [7] took advantage of the natural redundancy in a distributed shared memory system to make it resilient to single processor failures. Both of these systems export a shared-memory interface to the programmer and embed fault-tolerance into the implementation with no reliance on stable storage. Plank et al [26] embedded diskless checkpointing (with Raid Level 5 encoding) into several matrix operations in the ScaLAPACK distributed linear algebra package, thus making them resilient to single processor failures with low overhead. Kim et al [15] extended this work to employ one-dimensional parity encoding, which both lowers the overhead and increases the failure coverage. In [23], diskless checkpointing ideas are extended to a disk-based checkpointing system where there is disparity between the performance of local and remote disk storage. In such environments, diskless checkpointing may be extended so that in-memory checkpoints are stored on local disks (which are fast, but do not survive processor failures), and checkpoint encodings are stored on remote disks (which are slow, but are available following a failure). The performance of mirroring, Raid Level 5, and Reed-Solomon codings are all assessed and compare favorably to standard checkpointing to remote disk. The impact of checkpointing on the remote disk and the network is also assessed. Finally in [35], Vaidya makes the case for two-level recovery schemes, where a fast checkpointing method tolerating single processor failures is combined with a slower method that tolerates wholesale system failures. In his examples, checkpoint mirroring is employed for the fast method, and DISK-FORK checkpointing is employed for the slow method. His analysis applies to the methods presented in this paper as well. Diskless checkpointing is a technique where processor redundancy, memory redundancy and failure coverage are traded off so that a checkpointing system can operate in the absence of stable storage. In the process, the performance of checkpointing, as well as its impact on shared resources is improved. In this paper, we have described basic diskless checkpointing plus several performance optimizations. These have all been implemented and tested on five long-running application programs on a network of workstations and compared to standard disk-based checkpointing. In this implementation, the diskless checkpointing algorithms show a 34-fold improvement in checkpointing latency combined with comparable checkpoint overhead. The result is a lower expected running time in the presence of single processor failures. Several checkpointing systems [6, 23, 26, 32] have included variants of diskless checkpointing to improve the performance of checkpointing. Designers of checkpointing systems should consider the variants of diskless check-pointing presented in the paper to optimize performance and minimize the impact of checkpointing on shared resources. --R Virtual memory primitives for user programs. Application level fault tolerance in heterogeneous networks of workstations. EVENODD: An optimal scheme for tolerating double disk failures in RAID architectures. MIST: PVM with transparent migration and checkpointing. Efficient checkpoint mechanisms for massively parallel machines. Lightweight logging for lazy release consistent distributed shared memory. A survey of rollback-recovery protocols in message-passing systems The performance of consistent checkpointing. Manetho: Transparent rollback-recovery with low overhead A system for program debugging via reversible execution. Redundant Disk Arrays: Reliable Solutions to the shallow water test set using the spectral transform method. Fault tolerant matrix operations for networks of workstations using multiple checkpointing. Job and process recovery in a UNIX-based operating system Introduction to Parallel Computing. The checkpoint mechanism in KeyKOS. Low-latency, concurrent checkpointing for parallel programs. A longitudinal survey of internet host reliability. Improving the performance of coordinated checkpointers on networks of workstations using RAID techniques. A tutorial on Reed-Solomon coding for fault-tolerance in RAID-like systems Libckpt: Transparent checkpointing under unix. Fault tolerant matrix operations for networks of workstations using diskless checkpointing. Faster checkpointing with N Compressed differences: An algorithm for fast incremental check- pointing Transparent fault tolerance for parallel applications on networks of workstations. Checkpointing SPMD applications on transputer networks. Consistent checkpoints of PVM applications. The Condor distributed processing system. A case for two-level distributed recovery schemes Impact of checkpoint latency on overhead ratio of a checkpointing scheme. In 25th International Symposium on Fault-Tolerant Computing Demonic memory for process histories. --TR --CTR Sangho Yi , Junyoung Heo , Yookun Cho , Jiman Hong, Adaptive page-level incremental checkpointing based on expected recovery time, Proceedings of the 2006 ACM symposium on Applied computing, April 23-27, 2006, Dijon, France Kai Hwang , Hai Jin , Edward Chow , Cho-Li Wang , Zhiwei Xu, Designing SSI Clusters with Hierarchical Checkpointing and Single I/O Space, IEEE Concurrency, v.7 n.1, p.60-69, January 1999 Junyoung Heo , Sangho Yi , Yookun Cho , Jiman Hong , Sung Y. Shin, Space-efficient page-level incremental checkpointing, Proceedings of the 2005 ACM symposium on Applied computing, March 13-17, 2005, Santa Fe, New Mexico Xiaojuan Ren , Rudolf Eigenmann , Saurabh Bagchi, Failure-aware checkpointing in fine-grained cycle sharing systems, Proceedings of the 16th international symposium on High performance distributed computing, June 25-29, 2007, Monterey, California, USA Saurabh Agarwal , Rahul Garg , Meeta S. Gupta , Jose E. Moreira, Adaptive incremental checkpointing for massively parallel systems, Proceedings of the 18th annual international conference on Supercomputing, June 26-July 01, 2004, Malo, France Raphael Y. de Camargo , Renato Cerqueira , Fabio Kon, Strategies for storage of checkpointing data using non-dedicated repositories on Grid systems, Proceedings of the 3rd international workshop on Middleware for grid computing, p.1-6, November 28-December 02, 2005, Grenoble, France Ling , Jie Mi , Xiaola Lin, A Variational Calculus Approach to Optimal Checkpoint Placement, IEEE Transactions on Computers, v.50 n.7, p.699-708, July 2001 Adnan Agbaria , Hagit Attiya , Roy Friedman , Roman Vitenberg, Quantifying rollback propagation in distributed checkpointing, Journal of Parallel and Distributed Computing, v.64 n.3, p.370-384, March 2004 Daniel A. Reed , Charng-da Lu , Celso L. Mendes, Reliability challenges in large systems, Future Generation Computer Systems, v.22 n.3, p.293-302, February 2006 Zizhong Chen , Graham E. Fagg , Edgar Gabriel , Julien Langou , Thara Angskun , George Bosilca , Jack Dongarra, Fault tolerant high performance computing by a coding approach, Proceedings of the tenth ACM SIGPLAN symposium on Principles and practice of parallel programming, June 15-17, 2005, Chicago, IL, USA Milos Prvulovic , Zheng Zhang , Josep Torrellas, ReVive: cost-effective architectural support for rollback recovery in shared-memory multiprocessors, ACM SIGARCH Computer Architecture News, v.30 n.2, May 2002 Daniel J. Sorin , Milo M. K. Martin , Mark D. Hill , David A. Wood, SafetyNet: improving the availability of shared memory multiprocessors with global checkpoint/recovery, ACM SIGARCH Computer Architecture News, v.30 n.2, May 2002 Feng Qin , Joseph Tucek , Jagadeesan Sundaresan , Yuanyuan Zhou, Rx: treating bugs as allergies---a safe method to survive software failures, ACM SIGOPS Operating Systems Review, v.39 n.5, December 2005 Sudarshan M. Srinivasan , Srikanth Kandula , Christopher R. Andrews , Yuanyuan Zhou, Flashback: a lightweight extension for rollback and deterministic replay for software debugging, Proceedings of the USENIX Annual Technical Conference 2004 on USENIX Annual Technical Conference, p.3-3, June 27-July 02, 2004, Boston, MA

memory redundancy;checkpointing;RAID systems;error-correcting codes;rollback recovery;copy-on-write;fault tolerance

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Computing the Local Consensus of Trees.

The inference of consensus from a set of evolutionary trees is a fundamental problem in a number of fields such as biology and historical linguistics, and many models for inferring this consensus have been proposed. In this paper we present a model for deriving what we call a local consensus tree T from a set of trees ${\cal T}$. The model we propose presumes a function f, called a total local consensus function, which determines for every triple A of species, the form that the local consensus tree should take on A. We show that all local consensus trees, when they exist, can be constructed in polynomial time and that many fundamental problems can be solved in linear time. We also consider partial local consensus functions and study optimization problems under this model. We present linear time algorithms for several variations. Finally we point out that the local consensus approach ties together many previous approaches to constructing consensus trees.

Introduction An evolutionary tree (also called a phylogeny or phylogenetic tree) for a species set S is a rooted tree with leaves labeled by distinct elements in S. Because evolutionary history is difficult to determine (it is both computationally difficult as most optimization problems in this area are NP-hard, and scientifically difficult as well since a range of approaches appropriate to different types of data exist), a common approach to solving this problem is to apply many different algorithms to a given data set, or to different data sets representing the same species set, and then look for common elements from the set of trees which are returned. Several methods are described in the literature for deriving one tree from a set of trees. In this paper, we propose a new model, called the local consensus. This model is based upon functions, called local consensus rules, for inferring the rooted topology of the homeomorphic subtree induced by triples of species. We will show that any local consensus function can be Supported in part by NSF grant CCR-9108969. y Supported in part by ARO grant DAAL03-89-0031PRI z Supported in part by ARO grant DAAL03-89-0031PRI computed in polynomial time, and that many of the natural forms of the local consensus can be computed in linear time. We also analyze optimization problems based upon partial local consensus rules and show that many of these can also be solved in polynomial time. Preliminaries 2.1 Definitions Let be a set of species. An evolutionary tree for S (also known as a phylogenetic tree or, more simply, a phylogeny) is a rooted tree T with n leaves each labeled by a distinct element from S. The internal nodes denote ancestors of the species in S. For an arbitrary subset S 0 ae S we denote by TjS 0 the homeomorphic subtree of T induced by the leaves in S 0 . In particular, for a specified triple fa; b; cg ae S we denote by Tjfa; b; cg the homeomorphic subtree of T induced by the leaves labeled by a; b; and c. This topology is completely determined by specifying the pair of species among a; b and c whose least common ancestor (lca) lies furthest away from the root. If (a; b) is this pair then we denote this by ((a; b); c), and T is said to be resolved on the triple a; b; c. If T is not binary it may happen that all three pairs of species have the same least common ancestor. In this case we will say that is unresolved in T and denote this topology by (a; b; c). Given a tree T containing nodes u; v; w, we let lca T (u; v; w) denote the least common ancestor of u; v and w in T . Also, we let u - T v denote that v is on the path from u to the root of T . The set of input trees fT to a consensus problem is sometimes referred to as the profile. Let T (a; b; c) represent the set of rooted subtrees on the leaf set fa; b; cg. A local consensus rule is a function Given a local consensus rule f and a set R of evolutionary trees for S, the f-local consensus (if it exists) is a tree R f such that for all triples A ' S, R f When is said to be a total local consensus, and otherwise f is said to be a partial local consensus. The problem of determining if the f-local consensus exists and constructing it if it does is called the f-local consensus problem. We will also consider optimization versions of the local consensus problem which will be discussed in subsequent sections. Having set up this general machinery, we will look at the special case where we need to build a consensus of two trees and describe specific local consensus functions f for which we produce efficient algorithms. 2.2 Particular Local Consensus Rules We define the Binary Local Consensus, Optimistic Local Consensus and Pessimistic Local Consensus problems below. The Binary Local Consensus problem takes as input two binary trees, whereas the Optimistic Local Consensus and Pessimistic Local Consensus problems take as input two trees which are not necessarily binary. All of these are examples of total local consensus rules. Definition 2.1. A local consensus rule f is conservative if for every triple fa; b; cg for which T jfa; b; cg is required to be resolved for a particular profile, then no tree in the profile resolves fa; b; cg differently from T . When the trees are not necessarily binary, the local consensus rule has to interpret an unresolved triple in one of two distinct ways: supposing that any resolution of the three way split is possible or supposing that the unresolved node represents a three-way speciation event. Depending upon the interpretation, therefore, the local consensus rule may decide if T 1 is resolved and T 2 is unresolved on some triple, the output should be resolved identically to T 1 or unresolved. We call the first type of local consensus rule optimistic and the second type pessimistic. We now define these three consensus rules. Definition 2.2. Let T 1 and T 2 be two rooted binary trees on the same leaf set S. A rooted tree T (which is not necessarily binary) is called the binary local consensus of T 1 and T 2 iff for all triples a; b; c, T jfa; b; Definition 2.3. Let T 1 and T 2 be two rooted trees on the same leaf set S. A rooted tree T is called the optimistic local consensus of T 1 and T 2 iff for each triple a; b; c, c) or (a; b; c) for Definition 2.4. Let T 1 and T 2 be two rooted trees on the same leaf set S. A rooted tree T is called the pessimistic local consensus of T 1 and T 2 iff for each triple a; b; c, These differences are each appropriate for particular types of data. Given the above definitions of the three models, the local consensus tree may not exist. In Sections 3, 4, and 5, we will give linear time algorithms that either construct the tree we are looking for if it exists or conclude that no such tree exists. However, practicing biologists and linguists need to build some kind of consensus tree, and we therefore have considered variants of the local consensus tree problem which always have solutions. To this end, we will define the notion of relaxed-accord local consensus and relaxed-discord local consensus as follows. Definition 2.5. Let T 1 and T 2 be two rooted binary trees on the same leaf set S. A rooted tree T (which is not necessarily binary) is called a relaxed-accord local consensus of T 1 and whenever a triple a; b; c has differing topologies on T 1 and T 2 , that triple is unresolved in T and T preserves the topology of a maximal set of triples on which T 1 and T 2 agree. To prove the existence of a relaxed-accord local consensus tree it is sufficient to show that there exists a tree where every triple on which T 1 and T 2 disagree is unresolved. The set of trees with this property can be partially ordered based on the set of triples (on which T 1 and T 2 agree) whose topology they preserve. Once this partial order is known to be non-empty, we have proved the existence of a relaxed-accord local consensus since any maximal element in this partial order is such a consensus tree. We note that if T has the star topology it leaves unresolved all triples on which T 1 and T 2 disagree. Hence the partial order is non-empty and the relaxed-accord local consensus tree always exists. In Section 6 we show that this tree is unique. Definition 2.6. Let T 1 and T 2 be two rooted trees (not necessarily binary) on the same leaf set S. A rooted tree T is called a relaxed-discord local consensus of T 1 and T 2 if T preserves the topology of all triples on which T 1 and T 2 agree. In addition, T should leave unresolved a maximal set of triples on which T 1 and T 2 disagree. Using an argument similar to the one used to prove the existence of a relaxed-accord local consensus and noting that T 1 (or preserves the topology of all triples on which T 1 and agree, we conclude that the relaxed-discord local consensus always exists. In Section 6 we show that the relaxed-discord local consensus is also unique. Before we look further into the problems, we give some standard definitions available in the literature. Definition 2.7. Let T be a rooted tree with leaf set S. Given a node v 2 V (T ), we denote by L(T v ) the set of leaves in the subtree T v of T rooted at v. This is also called the cluster at v, and is represented by ff v . The set C(T is called the cluster encoding of T . Every rooted tree in which the leaves are labeled by S contains all singletons and the entire set S in C(T ); these clusters are called the trivial clusters. We define a maximal cluster to be the cluster defined by the child of the root. (Here we allow for a maximal cluster to be defined by a leaf also.) We also define the notion of compatibility of a set of clusters. Definition 2.8. A set A of clusters is said to be compatible iff there exists a tree T such that C(T The following proposition can be found in [12]. Proposition 2.1. A set A of clusters is compatible iff 8ff ;g. We now state a theorem which will be used in the later sections. Theorem 2.1. Let T 1 and T 2 be two rooted trees on the same leaf set S and let f be a conservative local consensus rule. If the f-local consensus tree T exists, then C(T are compatible sets. Proof. Suppose not and suppose without loss of generality that C(T ) [ not a compatible set. Then by Proposition 2.1, 9ff 2 C(T ) and fi 2 C(T 1 ) such that ff " ;g. Pick a ff. The topology of the triple a; b; c in T 1 is ((a; c); b) while in T it is ((a; b); c). Since f is a conservative local consensus rule, this is impossible. 2 3 Binary Local Consensus In this section, we will look at the Binary Local Consensus problem. We start by restating the definition of the binary local consensus tree : Let T 1 and T 2 be two rooted binary trees on the same leaf set S. A rooted tree T (which is not necessarily binary) is called the binary local consensus of T 1 and T 2 iff for all triples a; b; c, T jfa; b; 3.1 Characterization and construction We will show that althought the binary local consensus of two trees may not exist, when it does exist it has a nice characterization. Proposition 3.1. Given a binary tree T and a cluster ff, ff is compatible with C(T ) iff Proof. If ff 62 C(T ) but ff is compatible with C(T ), then there exists a proper refinement T 0 of T such that C(T 0 binary tree has a proper refinement. 2 Lemma 3.1. Let T 1 and T 2 be rooted binary trees on the same leaf set S. If f is a conservative local consensus function such that the f-local consensus tree T exists, then Proof. By the previous lemmas, C(T since f is conservative, and Corollary 3.1. If the binary local consensus tree T of T 1 and T 2 exists (T 1 and T 2 are both binary), then C(T Proof. All we need to show is that ff 2 C(T ) for all ff 2 C(T 1 For any such ff, pick identically, so that T also resolves The tree T such that C(T called the strict consensus tree. This particular consensus tree always exists and can be constructed in O(n) time [6]. The construction part of the algorithm for binary local consensus trees is therefore simple, and what remains is the verification that the strict consensus tree is also the binary local consensus tree (i.e. that the tree we have constructed using the algorithm in [6] satisfies the constraints imposed upon it by the binary local consensus rule). 3.2 Verifying that a Consensus Tree is a Binary Local Consensus Tree We now prove some structural lemmas to help determine whether the consensus tree is in fact the binary local consensus. Lemma 3.2. Let T 1 and T 2 be rooted binary trees on the same leaf set and let ff be a cluster in their intersection. Let T be the strict consensus tree of T 1 and T 2 . Let e 1 e be the edges in T respectively that are above the respective internal nodes which define the cluster ff. Let a be a species in ff. Then T is a binary local consensus for T 1 and T 2 if and only if 1. the subtree below e is a binary local consensus for the subtrees below e 1 and e 2 , and 2. upon replacing the subtrees below by a in T; respectively, T is a binary local consensus for T 1 and T 2 . Proof. Clearly, if T is the binary local consensus tree for T 1 and T 2 then conditions (1) and (2) will hold. Conversely, if (1) and (2) holds, but T is not the binary local consensus tree for T 1 and T 2 , then there is some triple a; b; c such that T incorrectly handles this triple. If all of a; b; c are below e then by condition (1), T handles a; b; c correctly. Similarly if at least two are above e, then by condition (2), T handles this triple correctly. It remains to show that T handles all triples where exactly two of a; b; c are below and one is above the edge e. But then, since the cluster in each of T this triple properly. Thus T is a binary local consensus for T 1 and T 2 . 2 This lemma yields an obvious divide-and-conquer strategy to determine whether a binary local consensus exists. Next we explore for what pairs of binary trees it is possible for the binary local consensus to be a star, i.e., a tree with none but the trivial clusters. Definition 3.1. A caterpillar is a rooted binary tree with only one pair of sibling leaves. Given a leaf labeled caterpillar T with root r and height h, there is a natural ordering induced by T on its leaves. Let hg be a function where g(s) is the distance of s from r. Then the species in S can be ordered in the increasing order as a 1 ; a such that g(a 1 (Note that the pair of sibling leaves have been arbitrarily ordered) Definition 3.2. Two caterpillars X and Y on the same leaf set are said to be oppositely oriented iff for all k, the k smallest elements of X are contained among the k+1 largest elements of Y and vice versa. Proposition 3.2. Let T 1 and T 2 be two rooted binary trees on the same leaf set whose binary local consensus is a star. If a; b is a sibling pair of leaves in T 1 , then the lca of a and b in T 2 must be the root of T 2 . Proof. Suppose not. Then there is a species c such that the least common ancestor of (a; c) is above the least common ancestor of (a; b) in T 2 . Then T 1 jfa; b; hence the binary local consensus of T 1 and T 2 cannot be a star. 2 Lemma 3.3. Suppose T 1 and T 2 are binary trees on the same leaf set and suppose that they each have at least 5 leaves. If their binary local consensus tree is a star, then T 1 and T 2 must be caterpillars. Proof. Suppose for contradiction that T 1 is not a caterpillar. Then it has two pairs of sibling leaves, (a; b) and (c; d). By the previous proposition each of these pairs must have the root as their least common ancestor in T 2 . Thus without loss of generality, a and c lie in the left subtree of the root of T 2 and b and d lie in the right subtree of the root of T 2 . Thus it follows that T 2 itself must have two sibling pairs (p; q) and (r; s) one in each subtree of the root. Note that in 1 the least common ancestor of p and q and the least common ancestor of r and s is the root of T 1 . Again without loss of generality let p and r lie in the left subtree of the root of T 1 and q and s lie in the right subtree of the root. x can be in either of the subtrees of T 1 Figure 1: Topologies of T 1 and T 2 with respect to p; q; Let x be any other species besides p; q; Figure 1). Suppose without loss of generality that x lies in the left subtree of the root of T 2 . We will consider the following two triples: In T 2 the topology of these triples will be ((x; p); s) and ((x; q); r) respectively. We will show that T 1 agrees on at least one of these triples. There are two cases. If x lies in the left subtree of the root of T 1 , then the topology of the triple x; p; s in T 1 is clearly ((x; p); s) and if x lies in the right subtree of the root of T 1 , then the topology of the triple x; q; r in T 1 is ((x; q); r). Thus in either case there is a triple in T 1 which agrees with a triple in T 2 and the binary local consensus cannot be a star. 2 Lemma 3.4. Let T 1 and T 2 be two caterpillars on the same leaf set. Then the binary local consensus of T 1 and T 2 is a star if and only T 1 and T 2 are oppositely oriented caterpillars. Proof. Suppose the two caterpillars are oppositely oriented, i.e., they satisfy the two intersection conditions. Let x; z be any three leaves and let their indices in the ordering of the leaves of T 1 be respectively. Then the topology of x; and z in T 1 is (x; (y; z)). Looking at the smallest elements in T 2 , this set must contain y or z but cannot contain x. Consequently, the topology of the triple in T 2 is not (x; (y; z)) and the star is a valid binary local consensus. Conversely, suppose that the two caterpillars do not satisfy the intersection conditions. Without loss of generality, suppose that there exists at least one k such that the k smallest elements of T 2 are not contained within the largest elements of T 1 . Pick the smallest such k. Say, x is the leaf in T 2 with rank k and x does not belong to the set of k largest elements of T 1 . From the pigeonhole principle, there will exist at least two leaves of T 2 which have ranks greater than k but which are contained in the set of k largest elements of T 1 . Suppose the two leaves are y and z. Then T 1 jfx; This implies that the binary local consensus cannot be a star. 2 a b c d e f d c a Figure 2: Example of oppositely oriented caterpillars Corollary 3.2. The binary local consensus for two trees can be verified to be a star in linear time. 3.3 Binary Local Consensus Tree Algorithm 1. Use Day's algorithm to produce the strict consensus tree T and for each non-trivial cluster in T , maintain a pointer to the edges in T 1 and T 2 that give rise to this cluster. 2. Traverse T in postorder. For each non-trivial cluster found, check that the subtrees below its edge in T 1 and T 2 are caterpillars that satisfy the conditions of the above lemma. If so replace the entire subtree by a single node belonging to the subtree in each of T 1 and T 2 . If not, declare that T is not the binary local consensus tree. Theorem 3.1. Construction and verification for the binary local consensus can be done in linear time. Proof. Day's algorithm [6] runs in linear time. Also, step 2 of the above algorithm takes linear time since at most a linear number of species are reintroduced by replacement above. Also, the checking of the caterpillars can be done in time linear in the number of leaves in the caterpillar.4 Optimistic Local Consensus In this section we look at the problem of finding the Optimistic Local Consensus (OLC) tree of two trees defined in the previous section. Note that the Optimistic Local Consensus of two trees may not exist. Recall the definition of the OLC tree : Let T 1 and T 2 be two rooted trees on the same leaf set S. A rooted tree T is called the optimistic local consensus of T 1 and T 2 iff for each triple c) or (a; b; c) for 4.1 Characterization of the OLC tree The following lemma characterizes the optimistic local consensus tree when it exists: Theorem 4.1. Let T 1 and T 2 be two rooted trees on the same species set S. If the optimistic local consensus tree T olc exists, then C(T olc and ff 2 2 C(T 2 ), and ff is compatible with both C(T 1 ) and C(T 2 )g. Proof. Pick any cluster ff 2 A. If we look at any triple x; then this triple will be resolved as ((x; y); z) in one tree and will be either resolved the same or unresolved in the other tree. In either case, ff 2 C(T olc ). Conversely, pick any cluster A. There are two cases here, namely, the case when ff is not compatible with at least one of C(T 1 ) and C(T 2 ) and the case when ff is compatible with both when ff is not compatible with at least one of C(T 1 ) and C(T 2 ), using Theorem 2.1, we observe that For the second case, pick those smallest clusters ff 1 . (Note that the nodes v and u defining the clusters ff 1 and ff 2 respectively, are the lcas in T 1 and T 2 respectively, of the species in ff.) Then 9fi ' ;, such that are the smallest clusters in T 1 and T 2 respectively containing ff and since ff is compatible with both C(T 1 ) and C(T 2 ), this implies that ff is the union of clusters of at least two children of v and also the union of clusters of at least two children of u. Moreover, 9a; b 2 ff such that (a; b) and (a; b). Thus we can pick a c 2 fi and we have that T 1 jfa; b; c). But the topology given by T olc is ((a; b); c). Thus 4.2 Construction phase Since the optimistic local consensus rule is conservative, if the tree set of clusters, and hence there exists a tree T satisfying C(T If we can construct T by refining T 1 , we can then reduce T by contracting all the unnecessary edges, and thus obtain T olc . This is the approach we will take. Note that this approach breaks the construction into two stages: refinement and contraction. Refining The main objective is to refine T 1 so as to include all the clusters from T olc . Before we explain how we do this precisely, we will introduce some notation and lemmas from previous works which enable us to do this efficiently. Definition 4.1. Let v be an arbitrary node in a tree T with children representative set of v is any set fx 1 such that x . We denote by rep(v) one such representative set. Lemma 4.1. If the optimistic local consensus tree T olc of trees T 1 and T 2 exists and then T olc jrep(v) is isomorphic to T 2 jrep(v). Proof. Follows from the fact that T 1 jrep(v) is a star. 2 Definition 4.2. Let v be a node in a tree T with children is the subtree induced by fv; g. We will do the refinement as follows. We will modify the tree T 1 is initialised to In a postorder fashion, for every It can be seen that v also has the same number of children as v (since the processing is done in a postorder fashion). Say these are . Replace the subtree (v ), rooted at v in the following manner : We replace N(v ) by an isomorphic copy of Next, we replace x i by the subtree of T 1 rooted at u i . Let T be the tree that is produced after considering all the nodes in T 1 . Theorem 4.2. Let T be given and suppose T olc exists. Then the tree T that is produced from the algorithm described in the previous paragraph satisfies C(T Proof. Since C(T olc all we need to show is that T olc jrep(v) cannot be a proper refinement of T 2 jrep(v). If it were, then for some fa; b; cg ' rep(v), T olc jfa; b; cg would be resolved while T 2 jfa; b; cg is unresolved. Since fa; b; cg ' rep(u), T 1 jfa; b; cg is also unresolved, forcing T olc to be also unresolved. 2 Note that we have reduced the problem of constructing T to the problem of discovering To have a linear time algorithm, however, we need to be able to compute T 2 jrep(v) quickly. We cite the following result from [13] which will be useful to us in this case. Lemma 4.2. [13] Given a left-to-right ordering of the leaves of a tree and the ability to determine the topology of any triple of leaves a; b; c in constant time, we can construct the tree in linear time. To use this lemma we need two things: 1. that we be able to determine the topology of any triple in T 2 in O(1) time, and 2. that we have for each node in T 1 , an ordered representative set, where the ordering is consistent with the left-to-right ordering of the leaves in T 2 . To accomplish (1), we first preprocess T 2 for lca queries. Then, to determine the topology for the triple a; b; c, we simply compare the lca's of (a; b), (b; c) and (a; c). The second requirement is more challenging, but can also be handled, as we now show. Computing all ordered representative sets in O(n) time: ffl Initially all nodes in T 1 have empty labelings. ffl For each s 2 S, taken in the left-to-right ordering of the leaves in T 2 , do 1. Trace a path in T 1 from the leaf for s towards the root, until encountering either the root or a node which has already been labeled. 2. Append s to the ordered set for each such node in the path traced (including the first node encountered which has already been labeled). Figure 3 shows an example of the computation just described. a b c d e a b c d e a b c d e c is added to rep sets of w and v Left-to-right ordering a c d b e a c d r a is added to rep sets of u, v and r (iv) After completion Figure 3: Example showing the computation of the representative sets of nodes in T 1 based on the left-to-right ordering of species in T 2 Note that this computation takes O(n) time since each node v is visited O(deg(v)) times, and that the order produced is exactly as required. Thus, for each node v 2 V (T 1 ), we have defined a set of leaves such that each leaf is in a different subtree of v, every subtree of v is represented, and the order in which these leaves appear is the same as the left-to-right ordering in T 2 . We have thus proved: Lemma 4.3. We can compute T 2 jrep(u) in O(jrep(u)j) time. We therefore have the following: Theorem 4.3. Given then we can construct a tree T such that C(T exists in O(n) time. The rest of the task of constructing T olc is in the contraction of unneeded edges. Contracting simply go through each edge in T and check if it needs to be kept or must be deleted. Note that edges that were added during the refinement phase are required and do not need to be checked. Therefore we need only check the original tree edges. Let (u; v) be such an edge with From our representative sets for u and v we can easily choose three species a; b; c such that lca(a; and lca(b; c) = v. If the topology of this triple in T 2 is differently resolved than ((a; b); c) then we know that edge (u; v) will have to be contracted; if on the other hand T 2 jfa; b; cg is either (a; b; c) or ((a; b); c) then (u; v) will have to be retained in any optimistic local consensus tree. OLC Construction Algorithm Phase 0: Preprocessing: Make copies T 0 2 of T 1 and T 2 respectively. For each node v in each tree T 0 compute ordered representative sets ordered by the left-to-right ordering in the other tree. Preprocess each tree T 0 i to answer lca queries for leaves as well as internal nodes. Phase I: Refine T 0Refine T 0 1 in a postorder fashion so that at the end, C(T 0 exists. Phase II: Contract T 0Contract edges e 2 E(T 0 1 ) such that c e , the cluster below e, lies in C(T 1 We have thus shown the following theorem Theorem 4.4. The algorithm stated above constructs the OLC of two trees T 1 and T 2 if the OLC exists. Analysis of Running Time Phase 0: Preprocessing: In [18], Harel and Tarjan give an O(n) time algorithm for preprocessing trees to answer lca queries in constant time. We have already shown that computing the ordered representative sets takes O(n) time. Thus the preprocessing stage takes O(n) time. Phase I: Refining T 0This stage involves local refinements of T 0 1 , and we have shown that the cost of refining around node v is O(deg(v)). Summing over all nodes v we obtain O(n) time. Phase II: Contracting edges This stage clearly takes only O(n) time. Theorem 4.5. Construction of the optimistic local consensus tree can be done in linear time. 4.3 Verification phase Lemma 4.4. Let T be a tree on a leaf set S. Let T be obtained from T through a sequence of refinements followed by a sequence of edge contractions. Then there exists a function there is a subset S v of the children of f(v) in Proof. We define set of clusters. Therefore there is a subset S v of the children of f(v) such that [ v 0 2Sv ff v Lemma 4.5. Suppose T is the OLC of T 1 and T 2 (on a leaf set S containing at least 5 species). Then T is a star iff either one of the following holds 1. both T 1 and T 2 are oppositely oriented caterpillars, or 2. both T 1 and T 2 are stars Proof. The "if" direction is easy to see. We now assume that the OLC, T , is a star. If contains a triple a; b; c that is unresolved, T 2 must also be unresolved on a; b; c. Conversely whenever T 1 is resolved on a; b; c, T 2 must be (differently) resolved on a; b; c. Thus either both are binary or both are not. In the case that both T 1 and T 2 are binary, the definition of the OLC coincides with the definition of binary local consensus and we appeal to the proofs of Lemma 3.3 and Lemma 3.4 to argue that T 1 and T 2 must be oppositely oriented caterpillars. are not binary, we will show that for any node v in T 1 with children fu there is a node v 0 in T 2 with children fu 0 k g such that ff u i . Pick any three species a; b; c such that a; b; c is unresolved in T 1 and let (a; b; c). Then a; b; c must be unresolved in T 2 . Let v (a; b; c). We claim that ff To see why, suppose being in the same subtree under v as a. Then This contradicts the assumption that T is a star. Thus ff Next, note that if x and y are under the same child of v in T 1 but under different children of there exists a z such that x; y; z is resolved in T 1 but unresolved in T 2 . This would contradict the fact the T is a star. This establishes the claim. This implies that if there is a non-binary node v that is not the root of T 1 , we can find two species a; b (a - v; b - v) and a species c, c 6- v such cg. Thus the root must have three or more children in this case. But this means that if any cluster defined by a child of the root contains two or more species, then there is a triple on which T 1 and T 2 agree. Thus T 1 and T 2 must be stars. 2 The verification proceeds as follows : Phase 0 Suppose the tree constructed by refining T 1 and then contracting the edges in the resulting tree is T . We will do the same modification on T 2 , i.e. refine T 2 using the information from T 1 and then contract the edges in the resulting tree as before. Call this tree T 0 . Clearly, if T is not isomorphic to T 0 , we can terminate and output that the OLC does not exist. This is because we know that a compatible set of clusters defines a unique tree and we know that the OLC, if it exists, is uniquely characterized. Phase 1 If Phase 0 is successful, we then verify further. We compute an ordered representative set for every node w in V (T ). For each node w in T , do 1. Check if the homeomorphic subtrees of T 1 and T 2 induced by rep(w) are both stars or they are both oppositely oriented caterpillars. If they are neither of these, then terminate and output that the OLC does not exist. 2. Identify the parent of w, say w . Look at rep(w ) excluding the representative element which is below w. Call this set A. Identify the lca's of rep(w) in T 1 and T 2 . Check if there is a species that belongs to A which lies below the lca of rep(w) in both T 1 or T 2 . If so, terminate and output that the OLC does not exist. Implementation of step 1 of Phase Using the left-to-right ordering of the species in T 1 , compute the ordered representative set rep, at each node in T as shown in the previous section. For any u 2 V (T ), to be able to quickly compute the homeomorphic subtree of T 2 induced by the species in rep(u), we need to know the ordering of theses species as they appear in the left-to-right ordering of T 2 . We associate with each u, a new rep set, rep (u), which is the rearranged version of the species in rep(u) according to their ordering in T 2 . We define a specifies for each s 2 S, the node v 2 V (T ) closest to the root of T such that s 2 rep(v). The function limit together with the left-to-right ordering of the species in T 2 , help in filling the rep sets, since, s will belong to the rep sets of all nodes in the path from s to limit(s). We first show how to compute using algorithm LIMIT and then we show how the rep sets are filled. Initialisation For each visited in a top-down traversal of T , do f For each s 2 rep(v) such that set Once limit(s) has been identified for all s 2 S, we proceed to compute rep (u); as follows. Look at the left-to-right ordering of the species in T 2 . Now, for each species s in the left-to-right order, we trace a path in T from the leaf for s towards the root of T and add s to the rep set of each node encountered in this path. We terminate when we reach limit(s). Note that this process of identifying rep and rep has to be done only once. Analysis of running time : The isomorphism test in Phase 0 can be performed in O(n) using a simple modification of the tree-isomorphism testing algorithm in [1]. There is an O(n) cost for preprocessing of T 1 and T 2 to answer lca queries in Phase 1. Our implementation of step 1 of Phase 1 involves a one time O(n) cost in preprocessing to identify rep and rep for each node in T . Then each time step 1 is called on a node w an additional time of O(deg(rep(w))) is taken. Exploiting that fact that T 1 and T 2 have been preprocessed to answer lca queries, it can be seen that each step 2 of Phase 1 takes O(deg(w) Thus the total time taken in the verification phase is O(n). Correctness of our verification procedure : Theorem 4.6. If T passes the above tests, then T is the OLC of T 1 and T 2 . Proof. We need only show that T handles every triple properly. Each of the following cases is handled assuming T has passed the isomorphism test. Case 1 If T passes the isomorphism test with T 0 , then any triple a; b; c such that the two trees resolve differently, will be unresolved in T . This follows since T is created by refining and then contracting both T 1 and T 2 , and these actions can not take a resolved triple into a different resolution. Case 2 This involves a triple a; b; c having the same topology ((a; b); c) in both T 1 and T 2 . We claim that the first step of Phase 1 will pass only if the topology of this triple is ((a; b); c). To see why, suppose a; b; c is unresolved in T . ( a; b; c cannot be resolved as (a; (b; c)) or ((a; c); b) in T .) Look at the nodes u and v, which are the lca's of a; b in T 1 and T 2 , respectively. The node w in T , which is the lca(a; b; c), is also lca(a; b) (since a; b; c is unresolved). We infer that f is the function as defined in Lemma 4.4. This is because, any node above w will contain the species c and any node below w will not contain either a or b. By a similar argument, Now, when we look at rep(w) and compute the homeomorphic subtrees of T 1 and T 2 induced by rep(w), in both of these induced trees, there will exist three species x; z such that x; y are both below u (and v) in T 1 (and T 2 ) and z is not in the character defined by u (and v). Thus in both the induced trees, the triple x; will have the same topology ((x; y); z). That is, these induced trees will neither be both stars nor both oppositely oriented caterpillars. Thus the verification process will terminate and output that the OLC does not exist. Case 3 This involves a triple a; b; c which is resolved as ((a; b); c) in one tree and unresolved in the other. The proof of this case essentially follows the lines of the proof of case 2. Case 4 This involves a triple a; b; c which is unresolved in both the trees. We claim that the second step of Phase 1 will pass only if this triple is unresolved in T . To see why, suppose a; b; c is resolved as ((a; b); c) in T . Let lca T (a; b; c) = x and let lca T (a; y and also suppose without loss of generality that x is the parent of y. Let y 1 be the child of y such that a 2 ff y 1 and let y 2 be the child of y such that b 2 ff y 2 . Let z 6= y be the child of x such that c 2 ff z . Let (a; b; c) and (a; b; c). We will look at functions f 1 and f 2 defined by Lemma 4.4 from V (T ) to V (T 1 respectively. v. Note that the cluster defined by any child of u can have a non-empty intersection with at most one of ff y 1 and ff y 2 . Similarly for v. Thus any representatives chosen from ff y 1 and ff y 2 respectively, have their least common ancestor at u in 1 and at v in T 2 . However, f 1 (z) - T 1 v. Thus any representative chosen from ff z will lie below u and v in T 1 and T 2 respectively, causing us to conclude that the OLC does not exist. 2 5 Pessimistic Local Consensus Recall the definition of the Pessimistic Local Consensus be two rooted trees on the same leaf set S. A rooted tree T is called the pessimistic local consensus of T 1 and T 2 iff for each triple a; b; c, T jfa; b; 5.1 Characterization The following theorem characterizes the PLC tree of two trees T 1 and Theorem 5.1. Let T 1 and T 2 be two trees on the same leaf set S. If the pessimistic local consensus tree T plc of T 1 and T 2 exists, then it is identically equal to T , where C(T Proof. Pick any cluster ff 2 C(T ). Since ff belongs to both the trees, if we look at any triple ff, then this triple will have to be resolved as ((x; y); z). Thus Conversely, pick any cluster We have two subcases here 1. ff is not compatible with at least one of C(T 1 ) or C(T 2 ). In this case, from Theorem 2.1, 2. ff is compatible with both C(T 1 ) and C(T 2 ). In this case, pick those nodes from T 1 and which define the smallest clusters containing ff. We can pick a triple a; b; c, such that a ff and this triple is unresolved in either T 1 or T 2 . Thus Construction Phase By Theorem 5.1, the pessimistic local consensus tree, if it exists, is identically the strict consensus tree. Thus to construct the pessimistic local consensus tree, it suffices to use the O(n) algorithm in [6] for the strict consensus tree. 5.3 Verification Phase Let T 1 and T 2 be the input trees, and let T be the strict consensus tree constructed using the algorithm in [6]. We want to be able to verify whether T is actually the pessimistic local consensus in the case that T is a star. If T 1 or T 2 is already a star then there is nothing to verify since T is the true pessimistic local consensus. So assume that this is not the case. There are two cases which we will consider. The first is when either of T 1 or T 2 (say T 1 ) has at least two children of the root which are not leaves. The second case is when both T 1 and T 2 have exactly one child of the root which is not a leaf. Having made observations about these cases, we can apply the divide and conquer strategy we adopted for the binary local consensus problem. Lemma 5.1. Suppose T 1 and T 2 are two trees on the same leaf set S, with T 1 having at least two children of the root which are not leaves. Let ff 1 ; ::; ff l be the maximal clusters of T 1 and be the maximal clusters of T 2 . Then T , their pessimistic local consensus, is a star iff Proof. Suppose 1. This means that 8x; y, if lca(x; y) in T 1 is below the root, then in T 2 , lca(x; y) is the root. Thus for any triple x; topologies in T 1 and T 2 do not agree. Thus T is a star. Suppose defined by a node which is not a leaf. Look at an ff k , such that the node in T 1 defining ff k is not a leaf node. There are two cases to handle here. Either, at least one species in ff k is not in fi j or all species in ff k are in fi j (i.e., ff k ae fi j ). In the former case, pick that species z, which is in ff k but not in fi j . Also pick those two species x; y which are in ff agree on the triple x; namely this triple has topology ((x; y); z) in both the trees. Thus T cannot be a star. In the latter case, since we know that fi j 6= S, we can pick two species x; y, from ff k and another species z, from S \Gamma fi j . In both T 1 and T 2 , the topology of this triple is ((x; y); z). Thus T cannot be a star. 2 Since each species belongs to at most one of these maximal clusters in each tree, this test can be done in linear time. The following lemma handles the case when both T 1 and T 2 have exactly one child of the root which is not a leaf. Lemma 5.2. Suppose T 1 and T 2 are two trees on the same leaf set S and T and their pessimistic local consensus, is a star. Suppose both T 1 and T 2 have exactly one child of the root each which is not a leaf. Let s be leaves in T 1 which are children of the root. Let v be the lca in T 2 of s . Then every child of v contains at most one species x g. Moreover, for any pair of species x; y g, the least common ancestor of x and y in T 2 lies on the path from v to the root. Proof. Suppose 9 a child of v which contains at least two species from S \Gamma fs g. Then by picking x; y such that they both lie under this child if v in T 2 and picking an s i out of s that lies under a different child of v, we find that both trees have the same topology for the triple cannot be a star. Furthermore, if 9x; y such that lca(x; y) in T 2 does not lie on the path from v to the root, then the triple x; would have identical topologies in both trees and T wouldn't be a star. 2 Definition 5.1. A rooted tree T is a millipede if the set of internal nodes of T defines a single path from the root to a leaf. a b c d Figure 4: An example of a millipede g. We have that T 2 jS 1 is a millipede (say, T Let l be the children of the root in T 2 , which are leaves. Look at T 1 jS 1 , (say, T Either, T 1 has one non-leaf child or it has at least two non-leaf children. In the former case, we can apply the previous lemma and infer that T will be a millipede. In the later case, we can apply Lemma 5.2 to check if the pessimistic local consensus is a star. In the following subsection, we will show how to verify if T is a star when both the input trees are millipedes. 5.3.1 Verification when both the input trees are millipedes The proof of the following lemma is straightforward. Lemma 5.3. Suppose T 1 and T 2 are two millipedes on the same leaf set S. Then their pessimistic local consensus, T , is a star iff there exists no triple such that both trees have the same topologies on the triple. We now describe a linear time algorithm for verifying that T 1 and T 2 have no triple on which they have the same topology. We define an ordering on the species in T 1 using the function f distance of s from the root of T 1 ; and, h is the height of T 1 . In T 2 , we can write S as the union of all the sets in the sequence S 1 is the height of T 2 and each S i contains exactly those species which are at a distance i from the root of T 2 . Now, in each S i , replace each species s in this set with f(s). Call this multiset of integers We thus get a sequence M of multisets. Definition 5.2. We will say a triple of integers p; q; r is special if We observe that the pessimistic local consensus of T 1 and T 2 is a star iff no special triple p; q and r exists. The following algorithm CHECK PLC, takes as input the sequence M returns FAIL if there exists a special triple of integers, and otherwise it returns PASS. CHECK PLC works by scanning the multiset M i in the i th iteration. It makes use of three variables global min, local min and temp. At the start of the i th iteration, global min stores the smallest integer seen in the first multisets. The variable local min is used to store the smallest integer a such that 9b for which a ! b and a 2 M (local min is initialised to +1). The variable temp is initialised to 0. As long as temp remains 0, local local min stores a and temp stores some b for which the previously mentioned realtionship between a and b holds. At the i th iteration, CHECK PLC either returns FAIL (if a special triple exists) or, if necessary, it modifies the variables global min, local min and temp to hold their intended values for the first i multisets of the sequence. The reasoning for storing these values at the start of the i th iteration is as follows. If 9p in some i) such that p; q; r is a special triple, then global min together with q; r 2 M i are also a special triple since global min - p. Similarly, if 9p in some M j , q 2 M l , i), such that p; q; r is a special triple, then local min, temp and r 2 M i are also a special triple. We now describe CHECK PLC. Initialisation: global local The procedure outputs FAIL (and terminates) if the pessimistic local consensus is not a star; it outputs PASS otherwise. For do f 2. do f Scan through M i ; If jAj - 2, then output FAIL; If y, where y 2 A local global If do f Scan through M i ; If either jAj - 2 or jBj - 1, then output FAIL; Else If If global min ! Min(M i ), then set local min = global min If global min ? Min(M i ), then set local min = global min y, where y 2 A global If Output PASS Analysis of running time : CHECK PLC runs in linear time since each M i is scanned only a constant number of times. Theorem 5.2. Algorithm CHECK PLC is correct. Proof. By induction, observe that Step 1 is executed at the i th iteration if 8j; l; x, where It then follows that if Step 1 is executed at the i th iteration, then at the start of that iteration local Thus, in this case global min stores the smallest integer seen in the first multisets. Now, in the first i multisets, if any special triple p; q; r exists such that and q; r 2 M i , then CHECK PLC correctly outputs FAIL since global min - p. Otherwise we have two cases, depending upon the value of A. If then the variables global min, temp and local min are updated so that global min holds the smallest value in the first i multisets. Also, local min, now correctly holds the smallest value a for which there exists a b (stored in temp) for which a ! b and a 2 M In the other case 0, in which case global min is updated to hold Min(M i ) (which is the smallest value in the first i multisets). Observe that once temp is updated to store a nonzero value, it never stores a 0 again. Thus, once temp is set to a nonzero value in iteration i 0 , then from iteration iteration k, Step 2 is executed. Assume that Step 2 is executed in some iteration i 0 and assume, inductively, that at the start of iteration i 0 , global min stores the smallest value in the first i multisets and local min stores the smallest value a for which there exists a b (stored in temp) such that a ! b and a 2 M . Then in iteration i 0 , it can be easily seen that CHECK PLC correctly outputs FAIL if there exist a special triple p; q; r such that both the cases when ensures that after iteration i 0 , global min stores the smallest value in the first i 0 multisets and local min stores the smallest value a for which there exists a b (stored in temp) such that a ! b and a 2 M Using the above arguments, it can be seen that CHECK PLC gives the correct output on any sequence of multisets. 2 Thus we also have the following theorem Theorem 5.3. Given two millipedes T 1 and T 2 , we can check if their pessimistic local consensus is a star in linear time. 6 Relaxed Versions The local consensus rules we have seen so far are such that the output tree satisfying a particular rule need not exist. This motivates the need to look at the relaxed versions of local consensus, where solutions always exist. Recall the definitions of relaxed-accord local consensus and relaxed- discord local consensus. The existance of solutions to these problems was shown in section 2.2. 6.1 Relaxed-Accord Local Consensus In this subsection we will show that the relaxed- accord local consensus of two binary rooted trees T 1 and T 2 is actually the strict consensus of these two trees. Theorem 6.1. If T 1 and T 2 are two rooted binary trees then their relaxed-accord local consensus T always exists, and is identically the strict consensus of T 1 and T 2 . Proof. The existence of the relaxed-accord local consensus tree T , was shown in section 2. Now we show that this tree is the strict consensus tree. Suppose there exists a triple a; b; c resolved differently in T 1 and T 2 , as say, ((a; b); c) and (a; (b; c)) respectively. Say the lca T 1 neither ff u nor ff v is in the strict consensus tree. Thus the strict consensus tree leaves unresolved any triple which has different topologies in T 1 and T 2 . Let T 0 be a tree in which for every triple a; b; c on which T 1 and T 2 differ, T 0 has an unresolved topology on this triple. Now suppose it is possible that T 0 contains a cluster that is not in the intersection of the sets of clusters of T 1 and T 2 . Let ff be this cluster and suppose without loss of generality that ff is not a cluster of T 1 . In T 0 , for any pair of species x; y 2 ff and species z 62 ff the topology has to be ((x; y); z). However, if this is also the case in T 1 , then T 1 must also possess the cluster ff contradicting our assumption. Thus there must exist a pair of species and a species z 62 ff such that in T 1 their topology is not ((x; y); z). But this implies that cannot be a relaxed-accord local consensus. Hence any candidate , T 0 , for a relaxed-accord local consensus can only contain the clusters in the intersection of the cluster sets of T 1 and T 2 . If T 0 contains a proper subset of the clusters in the intersection of the sets of clusters of T 1 and T 2 then there exists a triple a; b; c on which T 0 has an unresolved topology while the strict consensus tree has a resolved topology that agrees with the topologies of T 1 and T 2 . Hence the strict consensus of T 1 and T 2 is the relaxed-accord local consensus of T 1 and T 2 . 2 As a consequence, the relaxed-accord local consensus can be constructed in O(n) time using the algorithm in [6], and there is no need to verify that the tree constructed is correct. 6.2 Relaxed-Discord local consensus In the relaxed-discord local consensus (RDLC) problem we require that any triple on which the trees T 1 and T 2 agree must have its topology preserved in the consensus tree T . Further T should leave unresolved a maximal set of triples on which T 1 and T 2 disagree. Previously we showed that the RDLC exists. Now we will show that it is unique. The construction of the RDLC can be accomplished by defining the set b)c)g. This set of rooted triples can then be passed to the algorithm of Aho et al. [2], which computes a tree (if it exists) having the required form on every triple in the set, and resolving a minimum number of additional triples outside that set. The algorithm in takes O(pn) time where in our case, p 2 O(n 3 ), the use of the algorithm of would result in a running time of O(n 4 ). We will obtain a speed-up to an O(n 2 ) algorithm (which includes the verification) for the construction of the relaxed-discord tree, by using the fact that the tree necessarily exists. Our algorithm however takes advantage of the ideas in [2], and so we begin by briefly describing how that algorithm works. 6.2.1 The ASSU Algorithm In [2], Aho et al. describe algorithms which determine if a family of constraints on least common ancestor relations can be satisfied within a single rooted tree. We describe here the simple algorithm they give for the case where the constraints are given as rooted resolved triples, z). For such input the algorithm works top-down figuring out the clusters at the children of the root before recursing. To do this the algorithm maintains disjoint sets. Initially all leaves are in singleton sets. For each rooted triple ((x; y); z) the algorithm unions the sets containing x and y to indicate that x and y must lie below the same child of the root. This algorithm never unions sets unless this is forced. Recursive calls include constraints that are on species entirely contained in the same component discovered in the previous call. If all the species are seen to be in the same component (either initially or during a recursive call), the algorithm determines that the constraints cannot be simultaneously satisfied. This simple algorithm has a worst case behavior of O(pn), where there are p lca constraints and the underlying set S has n elements which will be leaves in the final tree. 6.2.2 An improved algorithm for RDLC We will now describe an algorithm to solve the RDLC in O(n 2 ) time. Since T 1 is itself consistent with all triples on which they agree, it is clear that T , the tree produced by the ASSU algorithm, is a refinement of this tree in the following sense. Each child of the root of T 1 (as well as represents a cluster which is the union of some of the clusters represented by children of the root of T . Let ff be the cluster of a child of the root of T 1 and let fi be a cluster of a child of the root of T 2 . ff and fi are unions of the clusters of some of the children of the root of T . In fact, if ff " fi is non-empty, then ff " fi is also the union of some of the clusters of the children of the root of T . We will show that except in one special case ff " fi is in fact the cluster of exactly one child of the root of T . fi). For any (y; z) in ff " fi, ((y; z); x) is the form of the triple within each of T 1 and T 2 and hence in T , y and z would lie under the same child of the root. Thus in this case ff " fi is a cluster of a child of the root of T . The case where ff [ can occur for at most one child of the root of T 1 and one child of the root of T 2 as the following lemma shows. Lemma 6.1. Let T 1 and T 2 be 2 trees on the same leaf set S. Let ff 1 ; ::; ff k be the clusters defined by the children of the root of T 1 and l be the clusters defined by the children of the root of T 2 . Then the case where ff i [ can occur for at most one i and one j. Proof. Suppose not. Let ff i [ with we have that ff i ' fi j . But since ff implies that This is a contradiction since fi j and fi j are clusters defined by the children of the root and hence should be disjoint. 2 The case for ff [ can be handled as follows. Identify the lca, say u, of the species in similarly, the lca, say v, of the species in S \Gamma fi in T 1 . Clearly, in T 2 , u will be a descendent of the node defining fi and in T 1 , v will be a descendent of the node defining ff. all the nodes in the path (in starting from the node defining fi, and ending at the node u. Let be the clusters defined by the children (not on the path) of the nodes in this path. Similarily, identify all the clusters defined by the children of the nodes in the path (in T 1 ), starting from the node defining ff, to the node v. Unioning pairs (x; y) whenever x and y lie in ff " fl i for some i or whenever they lie in we get a partition of components and these turn out to be exactly the clusters present at the children of the root of T . With the above characterization a high-level description of the algorithm to construct T can be given as follows: RDLC Construction Algorithm: 1. For each pair of maximal clusters ff 2 recursively compute the tree on ff " fi and make its root a child of the root of T . 2. If there are clusters ff and fi such that ff [ compute the partition of recursively compute the tree for each component of the partition and make the roots of these trees children of the root of T . Running Time Analysis: Note that this algorithm does not require an explicit verification of the constructed tree, since in fact we know that the tree exists and we are simply computing it by mimicking efficiently what the algorithm in [2] would create. There are at most n recursive stages. We will show that each stage can be implemented in proving the O(n 2 ) bound. Case 2 can be handled in O(n) time as follows. Build a graph with vertices labeled by species in ff " fi. Now for each i connect the vertices in ff " fl i by a path and do the same for each j and vertices in Find connected components of this graph in O(n) time. For each connected component Comp we will have to find homeomorphic subtrees of T 1 and T 2 whose leaf set is Comp and recurse on these subtrees. This task is common to both cases and is described after the discussion on Case 1. To handle Case 1 it is important not to waste time on empty intersections. So we consider each species in turn and label the intersection that this species lies in. Thus we will identify at most n non-empty intersections. Let ff " fi be one such intersection. We need to find a homeomorphic subtree of T 1 that has ff " fi as the leaf set. We will show how to do this in time proportional to the number of leaves in ff " fi. Assume that T 1 and T 2 have been preprocessed for least common ancestor queries. Also note that we know the left-to-right ordering of all leaves of T 1 as well as of T 2 . Given the leaves in left-to-right ordering is also known and is the one induced by the overall left-to-right ordering. By Lemma 4.2 we can reconstruct the topology of the tree in linear time. This is exactly what we need to show that one stage of the recursion can be accomplished in O(n) time and that the overall time for the algorithm is O(n 2 ). Clearly this case can be handled in linear time and can occur for at most one pair of children. 7 Polynomial Time Algorithms for Arbitrary Local Consensus Rules We show in this section some polynomial time algorithms for constructing local consensus trees. We begin by discussing the case where f is a partial local consensus function. Lemma 7.1. (Aho et. al[2]) Let A be a multi-set of k rooted triples on a leaf set S, with n. We can determine in O(kn log n) time if a tree T exists such that T jt is homeomorphic to t for all t 2 A. In [15], an algorithm is given for the problem addressed in [2] for the case where all the triples are resolved. In this case a faster algorithm can be obtained. Lemma 7.2. (Henzinger, King, Warnow [15]) Let A be a multi-set of k resolved rooted triples on a leaf set S, with n. We can determine in minfO(k whether a tree T exists such that T jfa; b; cg is homeomorphic to the rooted triple(s) in A on (if such a triple exists in A). Theorem 7.1. Let f be an arbitrary partial local consensus function and T a set of k evolutionary trees on S, with Then we can determine if the local consensus tree exists and construct it if it does in O(kn 3 ) time. Proof. Given f , T , and a triple A, we can determine the form of T f jA (for those triples A for which T f jA has a restricted form) in O(kn 3 ) time. By the previous lemma, we can determine if partial local consensus tree exists, and construct it if it does, in O(n 2:5 time. The total time is therefore bounded by the cost of computing the triples. 2 While partial local consensus trees can be constructed in O(kn 3 ), total local consensus trees can be computed even faster. Lemma 7.3. [Kannan, Lawler, Warnow [13]] Given an oracle O which can answer queries of "What is the form of T jfa; b; cg for a species set fa; b; cg?", we can construct in O(n 2 tree T consistent with all the oracle queries (if it exists), and O(rn log n) time if the tree T has degree bounded by r. Theorem 7.2. Let f be a total local consensus function. Then given a set of k rooted trees on n species, we can construct in O(kn 2 ) time the f-local consensus tree T f if it exists. If f always returns resolved subtrees, then we can compute T f in O(kn log n) time. Proof. We can implement the oracle determining the form of the homeomorphic subtree of T f on a triple a; b; c by first preprocessing the trees to answer least common ancestor (lca) queries in constant time, using [18]. Then, answering a query needs only O(k) time. By [13], we need only O(n 2 ) queries and O(n 2 ) additional work, for a total cost of O(kn 2 ) in the general case. When T f has degree bounded by r, we have total cost O(krn log n). If f always returns resolved subtrees, then T f will be binary, so that the total cost is O(kn log n). 2 8 Discussion and Conclusions Several approaches have been taken to handle the problem of resolving multiple solutions. One approach has been to find a maximum subset S 0 ' S inducing homeomorphic subtrees; this subtree is then called a Maximum Agreement Subtree[14, 10, 17]. The primary disadvantage of this approach is that it does not return an evolutionary tree on the entire species set. There is however a connection between this problem and one of the local consensus methods. The tree produced by the relaxed discord local consensus method contains the maximum agreement subtree as a homeomorphic subtree. This is not too hard to see. The other approach which we take here, requires that the resolution of the inconsistencies be represented in a single evolutionary tree for the entire species set. A classical problem in this area is the Tree Compatibility Problem (also called the Cladistic Character Compatibility Problem)[7, 8, 9] The Tree Compatibility Problem says that the set T of trees is compatible if a tree T exists such that for every triple A ' S, T resolves A if and only T which resolves A. This problem can be solved in linear time[12, 19]. The weakness of this approach is that in practice, many data sets are incompatible, and it is therefore necessary to be able to handle the case where some pairs of trees resolve triples differently. Some other approaches of this type are the strict consensus and the median tree problems. These models are stated in terms of unrooted trees, so that instead of clusters, characters (i.e. bipartitions) on the species set are used to represent the trees. Using the character encoding of the consensus tree as a measure of fitness to the input, the strict consensus seeks a tree with only those characters that appear in every tree in the input. The median tree, on the other hand, is defined by a metric, rooted trees which is defined to be the cardinality of the symmetric difference of the character sets of T 1 and T 2 . Given input trees T is the median tree if it minimizes P The median tree can be computed in polynomial time and has a nice characterization in terms of the character encoding [4, 16, 6]. Both the above notions are related to versions of the local consensus problem, and the relevant local consensus trees always contain at least as much 'information' as these trees. The work represented in this paper can be extended in several directions. As we have noted, for all local consensus functions the local consensus tree of a set of k trees can be computed in time polynomial in k and Many of these local consensus trees can be constructed in O(kn) time. --R The design and analysis of computer algorithms A formal theory of consensus The median procedure for n-Trees Mitochondrial DNA sequences of primates: tempo and mode of evolution Optimal algorithms for comparing trees with labeled leaves Optimal evolutionary tree comparison by sparse dynamic programming Numerical methods for inferring evolutionary trees Efficient algorithms for inferring evolutionary trees Determining the evolutionary tree Maximum agreement subtree in a set of evolutionary trees - metrics and efficient algorithms A fast algorithm for constructing rooted trees from constraints The Complexity of the Median Procedure for Binary Trees computing the maximum agreement subtree Fast Algorithm for Finding Nearest Common Ancestors --TR

graphs;algorithms;evolutionary trees

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First-Order Queries on Finite Structures Over the Reals.

We investigate properties of finite relational structures over the reals expressed by first-order sentences whose predicates are the relations of the structure plus arbitrary polynomial inequalities, and whose quantifiers can range over the whole set of reals. In constraint programming terminology, this corresponds to Boolean real polynomial constraint queries on finite structures. The fact that quantifiers range over all reals seems crucial; however, we observe that each sentence in the first-order theory of the reals can be evaluated by letting each quantifier range over only a finite set of real numbers without changing its truth value. Inspired by this observation, we then show that when all polynomials used are linear, each query can be expressed uniformly on all finite structures by a sentence of which the quantifiers range only over the finite domain of the structure. In other words, linear constraint programming on finite structures can be reduced to ordinary query evaluation as usual in finite model theory and databases. Moreover, if only "generic" queries are taken into consideration, we show that this can be reduced even further by proving that such queries can be expressed by sentences using as polynomial inequalities only those of the simple form x < y.

Introduction In this paper we are motivated by two fields of computer science which heavily rely on logic: relational databases and constraint programming. We will look at the latter from the perspective of the former. In classical relational database theory [1], a database is modeled as a relational structure. The domain of this structure is some fixed universe U of possible data elements (such as all strings, or all natural numbers), and is typically infinite. The relations of the structure, in contrast, are always finite as they model finite tables holding data. As a consequence, the active domain of the database, consisting of all data elements actually occurring in one or more of the relations, is finite as well. A (Boolean) query is a mapping from databases (over some fixed relational signature) to true or false. A basic way of expressing a query is by a first-order sentence over the relational sig- nature. For example, on a database containing information on children and hobbies, the query "does each parent have at least all hobbies his children have?" is expressed by the sentence (8p)(8c)(8h)(Child (p; c) - Hobby(c; Since the domain of each database is U, the quantifiers in a sentence expressing a query will naturally range over the whole infinite U. It is thus not entirely obvious that under this natural interpretation the query will always be effectively computable. That first-order queries are indeed computable follows immediately from a result by Aylamazyan, Gilula, Stolboushkin, and Schwartz [4] (for simplicity hereafter referred to as "the four Russians"). They showed that in order to obtain the result of the query it suffices to let the quantifiers range over the active domain augmented with a finite set of q additional data elements, where q is the number of quantified variables in the formula expressing the query. The intuition behind this result is that all data elements outside the active domain of a given database are alike with respect to that database. Alternatively, we can choose to let the quantifiers range over the active domain only, thus obtaining a semantics which is quite different from the natural interpretation. For example, consider databases over the single unary relation symbol . Then the sentence (8x)P (x) will always be false under the natural interpretation, while under the active-domain interpretation it will always be true. In fact, it is not obvious that each query expressible under the natural interpretation is also expressible under the active-domain interpreta- tion. Hull and Su [11] established that the implication indeed holds. (The converse implication holds as well, since the active-domain interpretation can easily be simulated under the natural interpretation using bounded quantification.) In recent years, much attention has been paid to "constraint programming languages" (e.g., [5]). In particular, in 1990, Kanellakis, Kuper and Revesz demonstrated that the idea of constraint programming also applies to database query languages by introducing the framework of "con- straint query languages" [12]. An important instance of this framework is that of real polynomial constraints. Here, the universe U of data elements is the field R of real numbers. Data-bases then are relational structures over R, but the database relations need no longer be finite; it suffices that they are definable as finite Boolean combinations of polynomial inequalities. In other words, each k-ary relation of the structure must be a semi-algebraic subset of R k [6]. A basic way of querying real polynomial constraint databases is again by first-order sentences, which can now contain polynomial inequalities in addition to the predicate symbols of the relational signature. For example, if the database holds a set S of points in R 2 , the query "do all points in S lie on a common circle?" is expressed by that quantifiers are naturally interpreted as ranging over the whole of R. In order to evaluate such a sentence on a database, we replace each predicate symbol in the formula by the polynomial definition of the corresponding data-base relation, and obtain a sentence in the pure first-order theory of the reals. As is well-known, this theory is decidable [15]; the truth value of the obtained sentence yields the result of the query. So, real polynomial constraint queries are effectively computable. Finite relations are semi-algebraic, so that finite relational databases over the reals form an important special case of real polynomial constraint databases. For example, if we want to model a database holding a finite number of rectangles, we can either choose to store the full extents of the rectangles, resulting in the infinite set of all points on the rectangles (represented in terms of linear inequalities in the obvious way), or we can choose to store only the corner points of each rectangle, resulting in a finite relation. In the present paper, we investigate whether the results by the four Russians and by Hull and Su, mentioned in the beginning of this Introduc- tion, carry over from classical first-order queries on relational databases to polynomial constraint queries on finite databases over the reals. In- deed, as in the classical case, one can give an alternative active-domain semantics to constraint sentences and again ask whether this is without loss of expressive power. Note, however, that active-domain quantification defies the very nature of constraint programming as a means to reason about intentionally defined, potentially in- finite, ranges of values. Hence, it is not obvious that the results just mentioned might carry over at all. Nonetheless, we have found a natural analog of the four Russians theorem, and we have been able to establish the verbatim analog of the Hull-Su theorem in the case when only linear polynomials are used. This is often the case in practice. Our result might be paraphrased by saying that on finite structures, first-order linear constraint programming can be reduced to ordinary query evaluation as usual in finite model theory and databases Our development is based upon the following observation. Consider a prenex normal form sentence in the first-order theory of the reals. For any finite set D 0 of real numbers, there exists a sequence D 0 ' D 1 ' of finite sets of reals such that the sentence can be evaluated by letting each quantifier range over D i only (rather than over the whole of R) without changing the sentence's truth value. By taking D 0 to be the active domain of a given finite database over the reals, we get the analog in the real case of the four Russians theorem. The reader familiar with Collins's method for quantifier elimination in real-closed fields through cylindrical algebraic decomposition (cad) [3, 8] will not be surprised by the above observation. In- deed, it follows more or less directly from an obvious adaptation of the cad construction. However, we give an alternative self-contained proof from first principles which abstracts away the purely algorithmical aspects of the cad construction and focuses on the logic behind it. Importantly, this proof provides us with a basis to show how in the case of linear polynomials, the construction of the sequence departing from the active domain D 0 can be simulated using a linear constraint formula. As a result, we obtain the analog in the real case of the Hull-Su theorem. In a final section of this paper, we look at queries that are "generic," i.e., that do not distinguish between isomorphic databases. Genericity is a natural criterion in the context of classical relational databases [2, 7]. Perhaps this is a little less so for databases over the reals; in other work [14] we have proposed alternative, "spatial" genericity criterions based on geometrical intu- itions. Nevertheless, it remains interesting to investigate which classically generic queries can be expressed using linear constraint sentences. Sentences that do not contain any polynomial inequalities always express generic queries, but from the moment a sentence even contains only simple inequalities of the form x ! y it can already be non-generic. Furthermore, there is an example due to Gurevich [1, Exercise 17.27], showing a generic query expressible with such simple inequalities but not without. In other words, simple inequalities, though inherently non-generic when viewed in isolation, help to express more generic queries. The natural question now is to ask whether general linear polynomial inequalities help even more. We will answer this question negatively, thus providing a partial rectification of Kuper's original intuitions [13] (which are incorrect as stated, by the Gurevich example just mentioned). Real formulas Let R be the field of real numbers. A real formula is a first-order formula built from atomic formulas of the form p ? 0, with p a multivariate polynomial with real coefficients, using logical connectives and quantifiers in the obvious manner. If are among x is a tuple of real numbers, then the satisfaction of \Phi on R with a i substituted for x i , denoted R defined in the standard way. As usual, a formula without free variables is called a sentence. Example 2.1 The formula \Phi(a; b; c) j (9x) a, b and c. [A condition like easily expressed in terms of conditions of the form p ? 0 by :(p ? 0)-:(\Gammap ? 0).] We have R but R 6j= \Phi[5; 3; 1]. be a real formula in prenex normal with each Q i either 9 or 8 and M quantifier-free. If D k+1 , . , D n are subsets of R, then we say that \Phi[-a] is satisfied on (D (D evaluates to true when we let each quantifier Q i range over D i only rather than over the whole of R. Example 2.2 Let \Phi be the sentence Our main result of this section can now be stated as follows: Theorem 2.3 Let \Phi(-x) be a prenex normal form real formula as in ( be a finite subset of R. Then there exists a sequence D k ' of finite subsets of R such that for all tuples - a on D k , R only if (D Example 2.4 As a trivial illustration let be the formula (9x 2 )x We have R we have (D 2 ) In the remainder of this section we give a simple proof of Theorem 2.3. We will introduce various auxiliary notions on which we will rely heavily in later sections. We first define the following natural equivalence relation on R Definition 2.5 Two points - a and - b in R n are called equivalent (with respect to \Phi), denoted - a j each polynomial occurring in \Phi has the same sign (positive, zero, or negative) on - a and - b. We now extend this equivalence relation inductively to lower dimensions such that the equivalence classes at each dimension are "cylindrical" over the equivalence classes at the next lower dimension Definition 2.6 Let i ! n and assume j is already defined on R i+1 . Then for have - a for each ff there is a fi such that conversely, for each fi there is an ff such that ( - b; fi) j (-a; ff) (with ff and fi real We note: Lemma 2.7 For each i, j is of finite index on R i . Proof. By downward induction on i. The case -a) the set of equivalence classes in R i+1 lying "above" - a, i.e., intersecting the line f(-a; ff) j ff 2 Rg. Clearly, so that we have an injection mapping each equivalence class c on R i to the set f-a) j - a 2 cg. Since j is of finite index on R i+1 , - can have only a finite number of possible values and hence j is of finite index on R i as well. The relevance of the equivalence relations just defined is demonstrated by the following lemma. We use the following stands for the formula (Q i+1 x M . Lemma 2.8 Let k - i - n, and let - a and - b be equivalent points in R i . Then R only if R This lemma can be proven by a straightforward induction (omitted). The notion of domain sequence, defined next, is crucial. The technical lemma following the definition will directly imply Theorem 2.3. Definition 2.9 Let D k be a finite subset of R. A sequence of finite subsets of R is called a domain sequence with respect to \Phi if for each k - Since j is of finite index, we know that a domain sequence always exists. Lemma 2.10 Let (D k ; D sequence w.r.t. \Phi, and let a only if (D Proof. By downward induction on i. The case n. Note that We concentrate on the case the case Q lar. Denote (a a. For the implication from left to right, assume R Then there exists a i+1 2 R such that R According to Definition 2.9, there exists a 0 By Lemma 2.8, since R also have R (D We can thus conclude that (D [-a]. The implication from right to left is straightforward. Theorem 2.3 immediately follows from the case of the above lemma. The reader may have noticed that we have never relied on the fact that in the polynomial inequalities p ? 0 occurring in a real formula, p is really a polynomial. So, the theorem holds for any first-order language of real functions. This observation substantiates our claim made in the Introduction that our proof "abstracts away the purely algorithmical aspects of Collins's cad construction and focuses on the logic behind it". Of course, by departing from the cad construction one gets an effectively computable version of Theorem 2.3. We will give an alternative construction for the linear case in Section 4. 3 Queries on real databases Fix a relational signature oe consisting of a finite number of relation symbols S with associated arity ff(S). A real database B is a structure of type oe with R as domain, assigning to each relation in oe a finite relation S B of rank ff(S) on R. The active domain of B, denoted by adom(B), is the (finite) set of all real numbers appearing in one or more relations in B. A query is a mapping from databases of type oe to true or false. A basic way of expressing queries is by first-order formulas which look like real for- mulas, with the important additional feature that they can use predicates of the form S(p where S is a relation symbol in oe of arity a, and each p i is a polynomial. If \Phi(-x) is a query formula and B is a database, then the satisfaction defined in the standard way. In par- ticular, if \Phi is a sentence, it expresses the query yielding true on an input database B iff B Example 3.1 Assume 2. The query "do all points in S lie on a common circle?" can be expressed as (9x 0 )(9y 0 )(9r)(8x) query "is there a point in S whose coordinates are greater than or equal to 1?" can be expressed as 1). Note that the quantifiers in query formulas are naturally interpreted as ranging over the whole of R. If \Phi is a query sentence and B is a database, then we can produce a real sentence \Phi B in a very natural way as follows. Let be an atomic subformula of \Phi, with S a relation symbol in oe. We know that S B is a finite relation consisting of, say, the m tuples )g. Then replace a ) in \Phi by It is obvious that B only if R Now assume the query sentence \Phi is in prenex normal If B is a database and D are subsets of R, then we say that \Phi is satisfied on (B; D evaluates to true on B when we let each quantifier Q i range over D i only, rather than over the whole of R. Given the preceding discussion, the following theorem follows readily from the material in the previous section: Theorem 3.2 Let \Phi be a query sentence as in (y) and let B be a database. If adom(B) ' D 1 ' is a domain sequence with respect to \Phi B , This theorem is the analog in the real case of the four Russians theorem [4] mentioned in the Introduction 4 The linear case In this section, we focus on linear queries, expressed by query sentences in which all occurring polynomials are linear. We prove that each linear query is expressible by a linear query sentence of which the quantifiers range over the active domain of the input database only. Thereto, we introduce a particular way to construct domain sequences on the active domain of a database, based on Gaussian elimination. We then show that this construction can be simulated in a uniform (i.e., database-independent) way by a linear query formula Before embarking, we point out that the notion of equivalence of points with respect to some given real formula \Phi (Definitions 2.5 and 2.6) depends only on the set of polynomials occurring in \Phi. So we can also talk of equivalence with respect to some given set of polynomials. Now let \Pi be a set of linear polynomials on the Such a polynomial p is of the form c p We define a sequence linear polynomials inductively as follows: c q In words, each \Pi i is a set of linear polynomials obtained from \Pi i+1 by Gaussian elimination. Equivalence of points in R i with respect to \Pi can be characterized in terms of the polynomials in \Pi i as follows: Proposition 4.1 Let 1 - i - n and let - Then - a j - b with respect to \Pi if and only if each polynomial in \Pi i has the same sign (positive, zero, or negative) on - a and - b. Proof. (Sketch) By downward induction on i. The case According to Definition 2.6, - a j - b if for each ff there is a fi such that (-a; ff) j ( - b; fi) (and conversely; for simplicity we will ignore this part in the present sketch). Equivalently, by induction, for each ff there is a fi such that each polynomial in \Pi i+1 has the same sign on (-a; ff) and ( - b; fi). For sim- plicity, we ignore in this sketch the polynomials equivalently, for each ff there is a fi such that for each p 2 \Pi i+1 , or ?. Now it can be seen that this is equivalent to p(-a)=c p for all By definition of \Pi i this is the same as saying that each p 2 \Pi i has the same sign on - a and - b. Now let \Phi be a linear query sentence (Q 1 prenex normal form, and let B be a database. Recall the definition of the real described in the previous section; note that since \Phi is linear, \Phi B is linear as well. Fix \Pi to be the set of polynomials occurring in \Phi B , and consider the sequence \Pi defined just above. We observe: Lemma 4.2 Let 1 - i - n. Then \Pi i is a finite union of sets of the form Both the number of these sets and the coefficients c i and d i for each set do not depend on the particular database B. Proof. Consider the case consists of the polynomials occurring in \Phi B . The elements of \Pi can be classified into two different kinds: those that already occur in \Phi, and those that are of the form e, with p a polynomial occurring in \Phi and e 2 adom(B). In the latter case, p\Gammae may be assumed to occur for all possible we omit the argument that this assumption is without loss of generality. Hence, the lemma holds for n. The case i ! n now follows easily by induction. We are now in a position to define a particular domain sequence with respect to \Phi B , based on the sequence Definition 4.3 The linear sequence on B with respect to \Phi is the sequence D 0 ' inductively defined as follows: D 0 equals adom(B), and for 1 where D 0 i is D Proposition 4.4 The linear sequence on B with respect to \Phi is a domain sequence with respect to Proof. According to Definition 2.9, we must show for each Consider the definition of D i in terms of D 0 4.3 above. We distinguish the following possibilities for ff: 1. 2. ff ? 3. is the maximal element in E i such that e 1 ! ff, and e 2 is the minimal element such that ff ! e 2 . It is obvious that ff 0 2 moreover, from the way defined, it is clear that all polynomials in \Pi i have the same sign on (-a; ff) and (-a; ff 0 ). Hence, by Proposition 4.1, the proposition follows. After one final lemma we will be able to state and prove the main result of this section: Lemma 4.5 For each 0 - i - n there exists a finite set P of linear polynomials such that for each database B, the i-th member D i of the linear sequence on B with respect to \Phi equals with z independent of B. Proof. By induction on i. The case is trivial since D The definition of D i in terms of D 0 i in Definition 4.3 is clearly of the consists of the four polynomials We have clearly of the form some P 00 , and by induction, D i\Gamma1 is of the form for some P 000 . By combining these expressions using a tedious but straightforward substitution process, we obtain the desired form for D i . Theorem 4.6 For each linear query sentence \Phi there is a linear query sentence \Psi such that for each database B, B \Psi, where adom denotes that the quantifiers in \Psi range over the active domain of the database only. Proof. Let be the linear sequence on B with respect to \Phi. By Theorem 3.2 and Proposition 4.4, we know that B We can write the latter explicitly as B case 8 is similar). From Lemma 4.5 we know that D 1 can be written as fp(y Pg. So, equivalently, we have where each (9y i ) ranges only over adom(B). By replacing in a similar manner, we obtain the desired sentence \Psi. 5 Generic queries Two databases B and B 0 over the same relational signature oe are called isomorphic if there is a for each relation symbol S in oe. A query which yields the same result on isomorphic databases is called generic. For example, assume that oe consists of a single binary relation symbol S. Databases of type oe can be viewed as finite directed graphs whose nodes are real numbers. Of course, any query expressed in the language L of pure first-order sentences over oe (i.e., not containing any polynomial inequalities) is generic. Other examples of generic queries are "is the graph connected?" or "is the number of edges even?". In the language L ! consisting of those query sentences where all inequalities are of the simple queries can be easily expressed, such as y. As pointed out in the Introduction, however, there are generic queries expressible in L ! but not in L. We have been able to prove that there is no similar gain in expressiveness when moving from L ! to full linear query sentences: Theorem 5.1 For each linear query sentence \Phi expressing a generic query there is a query sentence \Psi in L ! such that for each database B, As in Theorem 4.6, adom denotes that quantifiers range over the active domain only; we know by Theorem 4.6 that this active-domain interpretation is without loss of generality. Proof. (Sketch) We first observe: Lemma 5.2 Let p be a linear polynomial on the There exists a real formula involving only simple inequalities of the form x disjunction, conjunction, and negation, such that for each natural number s - 1, 1 As an aside, we would like the reader to note that it is possible to specialize Theorem 4.6 to query sentences in using a different construction of domain sequence. each sufficiently large real number a ? 0, and each tuple y on the set fa; a true. This lemma is proven by noting that the application of the (multivariate) polynomial p to can be viewed as the application of another, univariate polynomial to a. In particular, for a sufficiently large, the sign of the latter application is determined by the sign of the leading coefficient. The difficulty to be overcome is that this univariate polynomial depends on the particular y However, it can be seen that it depends essentially only on the way how the y are ordered. We omit the details Using the genericity of \Phi, we can now exploit the above lemma to prove the theorem as follows. Define \Psi to be the query sentence obtained by replacing each polynomial inequality p ? 0 occurring in \Phi by the corresponding formula / p as provided by the lemma. Now let B a database, and let s be the size of adom(B). Let a ? 0 be sufficiently large and let ae be an order-preserving bijection from adom(B) to fa; a g. Then we have \Psi. The first equivalence holds because \Phi is generic, the second equivalence is obvious from the lemma and the definition of \Psi, and the third equivalence holds because ae is order-preserving and \Psi 2 L ! (query sentences in L ! cannot distinguish between databases that are isomorphic via an order-preserving bijection). We can conclude that all generic queries that are not expressible in L ! , like even cardinality of a relation or connectivity of a graph, are not expressible as a linear query either. Non- expressibility in L ! has been addressed at length by Grumbach and Su [9]. Grumbach, Su, and Tollu [10] have also obtained inexpressibility results for linear queries, using complexity argu- ments. In particular, they showed that in the context of the rationals Q rather than the reals R, linear queries are the complexity class AC 0 , while even cardinality and connectivity are not. Acknowledgment We are grateful to Bart Kuijpers for his careful reading of earlier drafts of the material presented in this paper. --R Foundations of Databases. Universality of data retrieval languages. Geometric reasoning with logic and algebra. Constraint Logic Programming: Selected Re- search G'eometrie Alg'ebrique R'eelle. Computable queries for relational data bases. Quantifier elimination for real closed fields by cylindrical algebraic decom- position Finitely representable databases. Linear constraint databases. Domain independence and the relational calculus. Constraint query languages. On the expressive power of the relational calculus with arithmetic con- straints --TR --CTR Leonid Libkin, A collapse result for constraint queries over structures of small degree, Information Processing Letters, v.86 n.5, p.277-281, 15 June Gabriel M. Kuper , Jianwen Su, A representation independent language for planar spatial databases with Euclidean distance, Journal of Computer and System Sciences, v.73 n.6, p.845-874, September, 2007 Michael Benedikt , Martin Grohe , Leonid Libkin , Luc Segoufin, Reachability and connectivity queries in constraint databases, Proceedings of the nineteenth ACM SIGMOD-SIGACT-SIGART symposium on Principles of database systems, p.104-115, May 15-18, 2000, Dallas, Texas, United States Michael Benedikt , Martin Grohe , Leonid Libkin , Luc Segoufin, Reachability and connectivity queries in constraint databases, Journal of Computer and System Sciences, v.66 n.1, p.169-206, 01 February Michael Benedikt , Leonid Libkin, Relational queries over interpreted structures, Journal of the ACM (JACM), v.47 n.4, p.644-680, July 2000 Evgeny Dantsin , Thomas Eiter , Georg Gottlob , Andrei Voronkov, Complexity and expressive power of logic programming, ACM Computing Surveys (CSUR), v.33 n.3, p.374-425, September 2001

linear arithmetic;first-order logic;relational databases;constraint programming

291534

An Asymptotic-Induced Scheme for Nonstationary Transport Equations in the Diffusive Limit.

An asymptotic-induced scheme for nonstationary transport equations with the diffusion scaling is developed. The scheme works uniformly for all ranges of mean-free paths. It is based on the asymptotic analysis of the diffusion limit of the transport equation.A theoretical investigation of the behavior of the scheme in the diffusion limit is given and an approximation property is proven. Moreover, numerical results for different physical situations are shown and the uniform convergence of the scheme is established numerically.

Introduction . Transport equations are used to describe many physical phe- nomena. Some of the best known examples are neutron transport, radiative transfer equations, semiconductors or gas kinetics. The situation for small mean free paths is mathematically described by an asymptotic analysis. Depending on the transport equation and on the kind of scaling, different limit equations are obtained. For example the gas kinetic equations may lead to Euler or (in)compressible Navier Stokes equations. The limit equation for small mean free paths of radiative transfer, neutron transport, or semiconductor equations is the diffusion and the drift-diffusion equation, respectively. We refer to [3, 4, 12, 18, 20, 28] and [2, 6, 8, 10]. The main problem for numerical work on transport equations in these regimes is the stiffness of the equations for small mean free paths. For standard numerical schemes one has to use a very fine and expensive discretization with a discretization size depending on the mean free path. Moreover, in general a full resolution of the relaxation process is not necessary. The general aim is to develop numerical schemes working uniformly for different regimes. In particular, the discretization size should be independent of the mean free path. In recent years there has been a lot of work on numerical methods for kinetic equations in stiff regimes. For example, stationary transport equations in the diffusion limit have been considered, e.g., in [14, 15, 22, 21]. Nonstationary kinetic equations with a scaling leading to first order hydrodynamic equations like the Euler equation are treated in [7, 9]. Usually for the latter case a fractional step method with a semi-implicit treatment of the equations is used. For general work on implicit methods for transport equations we refer to [27] and references therein. We mention here also work on implicit methods for the full Boltzmann equation, see [5]. Moreover, the relaxation limit of transport equations may be used to develop schemes for the hydrodynamic equations themselves. These schemes have been developed by many authors. For a recent general approach to these so called relaxed or kinetic schemes we refer to [16]. The present work considers a scheme for nonstationary transport equations with a scaling leading to the diffusion equation as the limit equation. The different space time scalings involved in the problem are treated in a proper way. We use the standard perturbation procedure leading from the transport to the diffusion equation. FB Mathematik, University of Kaiserslautern, 67663 Kaiserslautern, Germany, ([email protected]). A. KLAR Essentially the problem is transformed into a system of equations of relaxation form and then a fractional step method is used. The analysis of the resulting problem is based on ideas developed in [7]. Including the results of a boundary layer analysis in the scheme, kinetic boundary layers are also treated in a correct way. Sections 2 and 3 contain a description of the results of the standard asymptotic procedure and the presentation of the time discretization in our scheme. In Section 4 the diffusion limit of the scheme is considered. In Section 5 the fully discretized equations are presented. An approximation property for different ranges of the mean free path is proven in 6. Section 7 contains numerical results for several examples and a numerical comparison with other schemes. Finishing the introduction we mention that the ideas developed in this paper can be transfered to the gas kinetic and the semiconductor case, where the above scaling leads in the limit to the incompressible Navier-Stokes equation and the drift-diffusion equation respectively. In particular, in the gas dynamic case a more careful use has to be made of the perturbation procedure leading from the Boltzmann equation to the incompressible Navier-Stokes equation. This problem will be treated in a separate paper. 2. The Equations. We consider transport equations of the following form where S is assumed to be the unit sphere around 0 in R 3 . The collision operator Q is defined by with are some constants and the scattering cross section oe is independent of v. K is an integral operator Z s symmetric in v and v 0 , rotationally invariant, are some constants, and R 1. K is compact. The collision operator has as collision invariants only constants and is negative in a suitable function space. The source term G(x) - 0 is assumed to be independent of v. Initial and boundary conditions are given by and where @D is the boundary of\Omega and the outer normal of @\Omega at the point x. See [3] for a thorough theoretical investigation of this equation. Extensions of the following to other cases like, e.g., v- dependence of oe and G are possible. Introducing the usual diffusion space-time scaling x ! x the mean free path and scaling G(x) one obtains the scaled equations ASYMPTOTIC-INDUCED SCHEME FOR TRANSPORT EQUATIONS 3 With the standard perturbation procedure, see e.g. [3, 4, 12, 20] the limit equation for (2.2) as ffl tends to 0 can be derived by writing f as One obtains and up to a constant be the solution of R 3. Since by assumption s is rotationally invariant, it follows [3] that 8i; j 2 R is the solution of the diffusion equation oe Doing a boundary layer analysis, one observes that the correct zeroth order boundary conditions for the diffusion equation are given by a kinetic half space problem: Let be the bounded solution of the following halfspace problem at x Then Here - x (\Gamma1; t) is independent of v. Remark: In the absorbing case the scaled equation (2.2) is changed into where oe A is the absorption cross section. The diffusion equation turns into oe 3. The Numerical Scheme. For a numerical scheme for the transport equation in the small mean free path limit it is desirable that varying mean free paths ffl can be treated with a fixed discretization such that it is not necessary to adapt the time step once the mean free paths tend to 0. Moreover, it is also desirable that the scheme is in the limit ffl ! 0 a good discretization of the diffusion equation. These points are obviously not fullfilled for a simple explicit time discretization of (2.2) like since, as ffl tends to 0, the time step must be shrinked due to stability considerations in order to treat the advection term (the CFL condition has to be fulfilled) and the 4 A. KLAR collision term properly. Therefore, large computation times are needed for small mean path for such a scheme. In contrast, for a fully implicit discretization there is no restriction on the time step due to stability considerations. However, one has to solve a stationary equation in every time step, which is again time consuming. We mention that, due to the development of fast multigrid algorithms [19, 24, 25, 26], for the stationary equation, computation times for a fully implicit scheme are strongly reduced. A numerical comparison of these types of algorithms with the one developed here is presented in Section 7. The aim in this work is to develop a semi-implicit scheme treating only such terms in an implicit way for which it is necessary to do so in order to obtain a scheme working uniformly in ffl. In particular, due to the different advection ( 1 scales, it is in the original formulation (2.2) not clear whether the advection has to be treated implicitely or not. One may nevertheless discretize the original equations in a straightforward way by treating the advection explicitely and the scattering term in an implicit way: This simple type of discretization has several drawbacks compared to the scheme developed below, we discuss them at the end of Section 4. We suggest to use the standard perturbation procedure to transform equation (2.2) into two equations. A fractional step scheme with a semi-implicit procedure is then used for the resulting equations. The idea is to follow the expansion procedure, collect suitable terms together, such that only terms on the scale 1 are involved. be the solution of the set of equations We take the inital and boundary values and One observes that f 0 fulfills the original equation (2.2) and the initial and boundary conditions. It is therefore the desired solution of the original problem. The results of the boundary layer analysis, see, e.g. [3], are included in the scheme by choosing h in the following way: be the solution of the halfspace problem (2.3). Since the outgoing ASYMPTOTIC-INDUCED SCHEME FOR TRANSPORT EQUATIONS 5 function at the boundary for the kinetic problem (2.2) is the same as the outgoing solution of the half space problem for ffl tending to 0, we define In the limit ffl tending to 0 we obtain in this way the correct boundary value. For independent of v we get q(x; v; It is obviously not reasonable to determine the outgoing function by solving the halfspace problem. This would need too much computing time. Here a fast approximate scheme as in [11] or [17] is needed to determine the outgoing function. For example a first approximation is given by choosing simply an approximation ~ of the asymptotic value of the halfspace problem as the outgoing function: The simplest approximation of - x (\Gamma1; t)is given by equalizing the half range fluxes of the halfspace problem at 0 and 1: ~ R R A more sophisticated approximation for q, see [17], is given by ~ R R (3.9)D4- Z (v R R dv and Z (w We remark that a correct treatment of the boundary conditions is important, in particular, if zeroth order kinetic boundary layers are present and one is using a coarse spatial grid not resolving the layer. See Section 7 for some examples. Using the approximations above one obtains a good approximation of the solution with a first order boundary layer even if only a very coarse grid is used. The first approximation yields in general already very good results as can be seen in the numerical experiments in Section 7. However, in certain situations the use of the second approximation might be necessary to obtain an improved accuracy, compare Figure 7.6 in Section 7. The system of equations (3.4,3.5) will be solved with a fractional step scheme: Step 2: 6 A. KLAR For Step 1 an explicit discretization will be used, Step 2 is discretized implicitely to treat the stiffness of the equations in a correct way. Let \Deltat denote the time step and f k \Deltat the time iterations approximating f 0 (x; v; k\Deltat); f 1 (x; v; k\Deltat). The initial and boundary values are given as above. Introducing the notation Z the time discretization is then given by the following: Step 2: \Deltat Rewriting (3.13,3.14) we obtain \Deltat \Deltat and \Deltat \Deltat This leads to Step 2: oe where the operator A is defined by \Deltatoe \Deltatoe I and \Deltat \Deltatoe I ASYMPTOTIC-INDUCED SCHEME FOR TRANSPORT EQUATIONS 7 Here I denotes the identity. The operator \Deltatoe I is positive and invertible with a bounded inverse in a suitable function space for all since K is compact and I \Gamma K is positive having only the constants as collision invariants. For example in the case of one-group transport with K =!? we obtain and In this case the semi-implicit scheme reduces to a fully explicit one. In general, in each time step we have to solve in Step 2 two linear Fredholm integral equation of the \Deltatoe I This may be achieved by standard methods [1, 13]. Remark: In the absorbing case one proceeds as described above treating the absorption in an explicit way. Then one obtains Step 2 is unchanged. 4. The Diffusion Limit. We start with the investigation of the behaviour of the time discretized scheme as ffl tends to 0 for fixed \Deltat. As the operators A and B have the following behaviour: Introducing suitable spaces, e.g. L 2 (S), we have for \Deltat fixed and ffl small, since positive the following: \Deltatoe and \Deltatoe Using these estimates we get that the scheme reduces in the diffusion limit, ffl tending to 0, to the following 8 A. KLAR Step 2: oe Moreover, we have and where h was defined in Section 2. This yields Step 2: oe Considering Step 2 and Step 1 together we obtain for oe oe or \Theta oe This is the simplest explicit time discretization for the diffusion equation. The boundary conditions for the diffusion equation that are given in the limit by the solution of the halfspace problem (2.3) fit to the boundary conditions for the kinetic scheme as defined in the last section. We finish this section by comparing the above scheme with the scheme (3.3) in Section 3. Doing the standard asymptotic analysis [21] we get for (3.3) as ffl ! 0 oe This means we obtain an explicit discretization of the diffusion equation as for the above scheme, but due to the ! f it is not the usual one. This type of discretization of the diffusion equation is worse in terms of accuracy and stability than (4.4). For example, doing a stability analysis one observes that only time steps are allowed which are half the size of those that can be used in (4.4). This is essentially due to the fully explicit treatment of the advection term in (3.3). Moreover, the scheme developed in Section 3 gives the possibility to treat for example the collision terms in a semi-implicit way as given in (3.13,3.14). This is at least for one-group transport with K =!? a decisive advantage, since the semi-implicit scheme presented here reduces in this case to a fully explicit one. If one would be trying to do the same thing based on the scheme (3.3) it would turn out that the limit equation is not any more the diffusion equation. ASYMPTOTIC-INDUCED SCHEME FOR TRANSPORT EQUATIONS 9 5. The Fully Discretized Equations. We restrict from now on for notational simplicity to the case, where f 0 and f 1 depend only on the first space coordinate: The domain under consideration is [0; L]. Moreover, we consider only the space discretiza- tion. The velocity space can be treated by using standard discretizations, see, e.g., [23]. We define a staggered grid x \Deltax , and x i\Gamma 1= We use the notation and Moreover the operators D + and D \Gamma are defined as The discretization of the initial values is straightforward. The boundary conditions are discretized by Discretizing the f 0 derivative in (3.16) with D \Gamma and the f 1 derivative in (3.11) with D+ yields the following scheme Step 2: oe A. KLAR In the limit for small ffl we obtain the space discretized diffusion equation oe \Deltat oe or \Theta \Deltat oe This is a standard explicit discretization of the diffusion equation. In particular, we obtain independent of the size of the discretization \Deltax a good discretization of the limit equation for all ranges of the mean free path. The discretization possesses all diffusion limits, the so called thin, intermediate and thick diffusion limit, see [22]. We observe, that we need in the limit a relation like \Deltat - (\Deltax) 2oe as for the diffusion equation, to obtain positivity and stability of our scheme. This condition may be relaxed for ffl large. 6. A Uniform Approximation Property. In this section we prove a uniform approximation property of our scheme. We give an estimate for the consistency error, considering the integral form of equations (3.4,3.5) assuming that the true solution is smooth. Written in integral form the equations for f 0 (t) and f 1 (t) are for the Cauchy problem and one space dimension ]ds: Approximating the integral by an integral over step functions defined in each interval of length \Deltat and approximating the derivative with respect to x as before, we get for oe oe where we defined f (k) ASYMPTOTIC-INDUCED SCHEME FOR TRANSPORT EQUATIONS 11 Reconsidering equations (5.1,5.2) and (5.3,5.4) for Step 1 and Step 2 of the numerical scheme and putting the steps together, we get for one time step or A or A oe oe +\Deltat A and a similar formula for f n 1 . In the following we want to estimate the consistency error, i.e. the difference between the true solution (f of the integral equation (6.1,6.2) and the value 1 ) that is obtained by introducing the true solution (f instead of (f k into the above formula (6.3). This means we estimate the difference between (f 0 (t); f 1 (t)) and ( ~ ~ A oe oe +\Deltat A and a similar formula for ~ 1 . Restricting in the proof for simplicity to the case I , we concentrate in the following on proving a pointwise estimate for The proof is based on four lemmas. Lemma 6.1. where C is a constant independent of ffl. Proof. A. KLAR Z (k+1)\Deltat k\Deltat oe oe oe ]ds oe oe Z (k+1)\Deltat k\Deltat oe \Deltax)ds oe oe k\Deltat oe ds where oe oe oe \Deltax oe used. Since the second term is 0, this is smaller than Lemma 6.2. A where C is a constant independent of ffl. Proof. A A with 1. This is equal to C \Deltat ASYMPTOTIC-INDUCED SCHEME FOR TRANSPORT EQUATIONS 13 since A. The estimates to prove the next Lemma can be found in [7] Lemma 6.3. with In particular for one has Proof. We have This gives \Deltat Moreover, since \Deltatoe \Deltatoe \Deltatoe we get \Deltatoe \Deltatoe \Deltatoe Estimates (6.4) and (6.5) give the first assertion. To prove the second assertion we use the first one with However, this is smaller than C \Deltat, since, if \Deltatoe 2 we have that it is smaller than if \Deltatoe that it is smaller than (noe \Deltat (noe \Deltat 14 A. KLAR Lemma 6.4. \Deltat where C is a constant independent of ffl. Proof. A oe oe The first two terms are estimated by Lemma 2 and Lemma 3. They are smaller than C \Deltat: The third term is for ffl - ffl 0 smaller than Using I the second term on the right hand side in (6.6) is smaller than \Deltatoe \Deltatoe \Deltatoe \Deltatoe \Deltat \Deltat Again due to Lemma 3 with k substituted by the first term in (6.6) is smaller than \Deltatoe \Deltatoe \Deltatoe \Deltatoe \Deltatoe ASYMPTOTIC-INDUCED SCHEME FOR TRANSPORT EQUATIONS 15 since x). For \Deltatoe smaller than For \Deltatoe smaller than \Deltatoe \Deltatoe \Deltatoe \Deltatoe \Deltatoe due to 1. This means that the first term in (6.6) is smaller than \Deltat \Deltat with C independent of ffl. Collecting all the terms the lemma is proven. All together, using Lemma 1 and 4 we have proven for t 2 [0; T ]; \Deltat This means, that for small \Deltat; \Deltax and fixed ffl, the estimate tends to 0 like \Deltat \Deltax. However, also for a meshsize, that is large compared to ffl the estimate shows, that we get convergence to 0. For example, for ffl - C \Deltat we obtain convergence to 0 like \Deltat \Deltax. We mention that \Deltat has to be chosen in relation to \Deltax . E.g. in the diffusion limit we need \Deltat to be of the order of (\Deltax) 2 as we have seen in the last section. 7. Numerical Results and Examples. In this section a numerical study of the scheme is presented and the scheme is compared with fully explicit and fully schemes. We restrict to the one-group transport equation in slab geometry, i.e. x 2 [0; L] and K =!?. This yields 3 . The velocity discretization is done using in all situations a 16 point quadrature set. We compute the solution with the semi-implicit scheme derived above for different space discretizations. To obtain positivity and stability of the semi-implicit scheme in the limit ffl tending to 0 one has to take - for a fixed space discretization \Deltax - a time step \Deltat of the size given by (5.6). As mentioned above this can be relaxed for large ffl. In particular, this means that the size of \Deltat can be chosen independent of ffl. Comparison with the explicit scheme (3.1): In contrast to the above we get that the explicit discretization (3.1) of equation (2.2) requires a time step of the order min(\Deltaxffl; ffl 2 oe ) to obtain positivity and stability. In particular, for small ffl the step size \Deltat has to be chosen in this case of the order ffl 2 , in contrast to the semi-implicit scheme. A comparison of the CPU time necessary for one time step yields that the semi-implicit scheme needs about 2 times the CPU time of the explicit scheme. This yields a big gain in computing time for small ffl for the semi-implicit scheme compared to an explicit one. In particular, it is reasonable to use A. KLAR the semi-implicit scheme, if 2 \Delta oe 2 oe and if the desired accuracy does not require a smaller time step, than the one that can be taken for the semi-implicit scheme. To obtain a certain required accuracy of the solution one has to use time steps as shown in the table below for some examples, see Table 1. Looking at Table one observes that using an explicit scheme is not reasonable for small ffl. Either the semi-implicit or the implicit scheme are faster. However, this changes for large, where the explicit scheme may be better due to the small computation times per time step. Comparison with the fully implicit scheme (3.2): A fully implicit dicretization of the equation obviously allows bigger time steps, since there is no stability restriction on the time step in this case. Nevertheless, for the accurate simulation of the time development small time steps may be necessary. To get an accurate resolution of the behaviour of the solution up to an error of a certain order the size of the time step for the implicit scheme has to be chosen according to Table An implementation of a fully implicit scheme shows that in order to obtain a sufficient accuracy the stationary equation has to be evaluated to a very high accuracy approximately up to an error of the order 10 \Gamma8 . To achieve this a standard iteration scheme using for example a diamond difference discretization needs a large number of iteration steps (sweeps over the computational domain). A comparison of the CPU time for one iteration step shows that one time step of the semi-implicit iteration needs less than 2 times the CPU-time of an iteration of the stationary scheme. Table 2 shows that the semi-implicit scheme has a big advantage compared to a standard implicit iteration in many situations. However, of course, computation times for an implicit scheme are strongly reduced if a multigrid algorithms as described, e.g., in [24] is used. Using the convergence estimates in [24] one observes that in essentially two cycles an accuracy of the one needed for the solution of the stationary equation is obtained. One V (1; 1) cycle costs about the same CPU time as 4 sweeps over the computational domain. I.e. the estimated costs for one time step of a fully implicit scheme with a multigrid algorithm is about 4 times as large as the one for the semi-implicit scheme. The complexity of the implementation of a multigrid scheme especially in higher dimensions has to be taken into consideration as well. We consider a situation with conditions equal to 0 at and equal to 1 at 1. The space discretizations are 0:01. The time steps required for the semi-implicit scheme by stability considerations are in this case \Deltat = 0:015; \Deltat = 0:00015. We consider final times 0:5. For a stationary state is nearly reached. The error was calculated by taking the L 1 -norm of the difference with the 'true ' solution computed with a very fine discretization. The table shows the time steps necessary to obtain a certain accuracy e with the implicit scheme using a diamond difference discretization. Table 1: Time steps required to obtain a certain accuracy e. These accuracy requirements together with the above estimated CPU time give the following relation between the CPU time for the explicit (E), the semi-implicit (S) ASYMPTOTIC-INDUCED SCHEME FOR TRANSPORT EQUATIONS 170.00050.00150.0025 Fig. 7.1. Error for different values of ffl. and the implicit scheme with multigrid (IM) and with a standard iteration procedure Table 2: Relative CPU times This shows that for coarse grids the semi-implicit scheme has to be prefered. For finer grids and nearly stationary situations the advantage of a fully implicit scheme with multigrid is clearly seen. Implicit schemes with a standard iteration procedure are in all considered situations slower than the semi-implicit scheme. Further investigation of the semi-implicit scheme: To show the uniform convergence in ffl for the semi-implicit scheme numerically, we compute the error for different values of ffl ranging from As before, we use conditions equal to 0 and 1 at and respectively, and the following values for the space discretization \Deltax with the corresponding \Deltat values due to the stability condition (5.6): Hence, each cell contains between 0:125 and 10 5 mean free paths. The error was calculated by taking the L 1 -norm of the difference of the solutions with discretization size 0.0125 and 0.025 (error 3), 0.025 and tively. This results in three curves, which are plotted in Figure 7.1. Looking at the figure, one observes that the error behaves perfectly uniform as ffl ! 0. The solution of the kinetic equation computed by the new scheme derived in this work is in the following computed for different physical situations. The physical examples under consideration are Example 1: the boundary conditions Example 2: the boundary conditions are equal to 0 and A. KLAR0.30.50.70 x semi-implicit25 true Fig. 7.2. Example 3: conditions equal to 0, Example 4: boundary conditons 0:4. The solution of this problem has a kinetic boundary layer at Example 5: As Example 4, but with Example A two material problem. In [0; 0:1] we consider a purely absorbing material with 0:1. I.e. the region has the size of one mean free path. In [0:1; 1:1] we take a purely scattering material region). The boundary conditions are f(0; The solution of this problem has an interface layer at The initial condition is always 0. The solutions for the physical situations described above are plotted in the following figures. In Figure 7.2 to 7.4 the situations from Example 1 to 3 are shown. The solutions are plotted using space discretizations for the semi-implicit scheme. We use the label 'semi-implicit10' to denote the solution with the semi-implicit scheme with 10 spatial cells. The time discretization is chosen due to the stability condition (5.6) for Example 2 and 3. For Example 1 the restriction on the time step is relaxed to a CFL-type condition. The reference solution is the solution with a very fine discretization. For this case the solution of the semi-implicit scheme and of the other schemes are coincident. The solution of the diffusion equation is computed by the usual triangular explicit scheme, which is the limiting scheme of our semi-implicit scheme as ffl tends to 0, compare (4.3). The example shows that for isotropic boundary conditions the solution is approximated with good accuracy for different ranges of ffl. In Figure figure layer1 Example 4 is considered. We plot the reference solution and the solution of the diffusion equation with boundary coefficients derived from the halfspace problem. The solutions of the semi-implicit scheme are found with such that a discretization cell contains and the corresponding size of the time discretization. The boundary values are found by determining approximately the outgoing distribution of the halfspace problem (2.3) as described in Section 3. ASYMPTOTIC-INDUCED SCHEME FOR TRANSPORT EQUATIONS 190.050.150.250 x true diffusion Fig. 7.3. x semi-implicit25 true Fig. 7.4. This is done using first the approximation of the asymptotic value of the halfspace problem by (3.7), as the outgoing function (the solution in the plot is labeled 'semi- implicit10-1') and second an outgoing function determined by formula (3.10) labeled 'semi-implicit10-2'. In this first case the two approaches give coincident results. One observes that even for a coarse diffusive discretization the behaviour of the solution at the boundary is found with very good accuracy. We mention that other approaches to obtain the correct discrete boundary conditions for the stationary equation are shown in [14, 21]. Figure 7.6 shows Example 5. The same as in Figure 7.5 is shown. However, in this case one cell contains now 1000 mean free paths. The advantage of using here an exact approximation of the outgoing fiunction of the half space problem is clearly seen. Finally Figure 7.7 shows Example 6. The space discretization is here in the absorbing region and \Deltax = 0:1 in the scattering region. In particular, one cell in the scattering region contains 100 mean free paths. The situation at the interface A. KLAR0.250.350.450.550.650.75 x true semi-implicit10-2 diffusion Fig. 7.5. x true semi-implicit10-2 Fig. 7.6. is treated in the same way as the one at the boundaries before. One observes again a good agreement of the solution in the diffusive region with the true solution. We mention here the work of [14, 18, 21] who treated similar problems for the stationary equation. ASYMPTOTIC-INDUCED SCHEME FOR TRANSPORT EQUATIONS 210.020.060.10.14 x true Fig. 7.7. 8. Conclusions. From the analytical and numerical results one can conclude: ffl The semi-implicit scheme works uniformly for all ranges of the mean free path. This is shown by numerical experiments and a consistency proof. ffl The limiting scheme for small mean free paths is a standard explicit discretization of the diffusion equation. ffl By including a boundary layer analysis one obtains a suitable treatment of the boundary conditions for coarse (diffusive) discretizations. ffl A comparison of the scheme with fully explicit and fully implicit schemes shows advantages and disadvantages. In particular, the semi-implicit scheme is faster than the fully implicit scheme, if the detailed time development is computed with a coarse discretization or with higher accuracy requirements. However, for nearly stationary situations with a fine grid the fully implicit scheme, if combined with a fast multigrid method as in [24], is faster. ffl The numerical results have been generated for the one group transport case. A further numerical treatment should include the implementation of the scheme with other scattering ratios. Using methods as in [1, 13] this should be possible without too much difficulties. --R A Survey of Numerical Methods for the solution of Fredholm Integral Equations of the Second Kind Fluid dynamic limits of kinetic equations: Formal derivations Diffusion approximation and computation of the critical size Boundary layers and homogenization of transport processes Implicit and iterative methods for the Boltzmann equation The fluid dynamical limit of the nonlinear Boltzmann equation Uniformly accurate schemes for hyperbolic systems with relaxation The Boltzmann Equation and its Applications Numerical passage from kinetic to fluid equations Incompressible Navier Stokes and Euler limits of the Boltzmann equation A numerical method for computing asymptotic states and outgoing distributions for kinetic linear half space problems Uniform asymptotic expansions in transport theory with small mean free paths Theorie und Numerik The relaxation schemes for systems of conservation laws in arbitrary space dimensions Diffusion theory as an asymptotic limit of transport theory for nearly critical systems with small mean free paths Asymptotic solution of neutron transport problems for small mean free path Asymptotic solution of numerical transport problems in optically thick Asymptotic solution of numerical transport problems in optically thick Computational Methods of Neutron Transport A fast multigrid algorithm for isotropic transport problems I: Pure scattering A fast multigrid algorithm for isotropic transport problems II: With absorption Multilevel methods for transport equations in diffusive regimes Methods of Numerical Mathematics Diffusion approximation of the linear semiconductor equation --TR

asymptotic analysis;diffusion limit;transport equations;numerical methods for stiff equations

291683

On Residue Symbols and the Mullineux Conjecture.

This paper is concerned with properties of the Mullineux map, which plays a rle in p-modular representation theory of symmetric groups. We introduce the residue symbol for a p-regular partitions, a variation of the Mullineux symbol, which makes the detection and removal of good nodes (as introduced by Kleshchev) in the partition easy to describe. Applications of this idea include a short proof of the combinatorial conjecture to which the Mullineux conjecture had been reduced by Kleshchev.

Introduction It is a well known fact that for a given prime p the p-modular irreducible representations D - of the symmetric group S n of degree n are labelled in a canonical way by the p-regular partitions - of n. When the modular irreducible representation D - of S n is tensored by the sign representation we get a new modular irreducible representation D - P . The question about the connection between the p-regular partitions - and - P was answered in 1995 by the proof of the so-called "Mullineux Conjecture". The importance of this result lies in the fact that it provides information about the decomposition numbers of symmetric groups of a completely different kind than was previously available. Also it is a starting point for investigations on the modular irreducible representations of the alternating groups. From a combinatorial point of view the Mullineux map gives a p-analogue of the conjugation map on partitions. The analysis of its fixed points has led to some interesting general partition identities [AO], [B]. The origin of this conjecture was a paper by G. Mullineux [M1] from 1979, when he defined a bijective involutory map - M on the set of p-regular partitions and conjectured that this map coincides with the map - P . The statement is the Mullineux conjecture. To each p-regular partition Mullineux associated a double array of integers, known now as the Mullineux symbol and the Mullineux map is defined as an operation on these symbols. The Mullineux symbol may be seen as a p-analogue of the Frobenius symbol for partitions. Before the proof of the Mullineux conjecture many pieces of evidence for it had been found, both of a combinatorial as well as of representation theoretical na- ture. The breakthrough was a series of papers by A. Kleshchev [K1], [K2], [K3] on "modular branching", i.e. on the restrictions of modular irreducible representations from S n to S n\Gamma1 . Using these results Kleshchev [K3] reduced the Mullineux conjecture to a purely combinatorial statement about the compatibility of the Mullineux map with the removal of "good nodes" (see below). A long and complicated proof of this combinatorial statement was then given in a paper by Ford and Kleshchev [FK]. In his work on modular branching Kleshchev introduced two important notions, normal and good nodes in p-regular partitions. Their importance has been stressed even further in recent work of Kleshchev [K4] on modular restriction. Also these notions occur in the work of Lascoux, Leclerc and Thibon on Hecke Algebras at roots of unity and crystal bases of quantum affine algebras [LLT]; it was discovered that Kleshchev's p-good branching graph on p-regular partitions is exactly the crystal graph of the basic module of the quantized affine Lie algebra U q ( - studied by Misra and Miwa [MM]. From the above it is clear that a better understanding of the Mullineux symbols is desirable including their relation to the existence of good and normal nodes in the corresponding partition. In the present paper this relation will be explained explicitly. We introduce a variation of the Mullineux symbol called the residue symbol for p-regular partitions. In terms of these the detection of good nodes is easy and the removal of good nodes has a very simple effect on the residue symbol. In particular this implies a shorter and much more transparent proof of the combinatorial part of the Mullineux conjecture with additional insights (Section 4). We also note that the good behaviour of the residue symbols with respect to removal of good nodes allows to give an alternative description of the p-good branching graph, and thus of the crystal graph mentioned above. Some further illustrations of the usefulness of residue symbols are given in Section 3. This includes combinatorial results on the fixed points of the Mullineux map. Basic definitions and preliminaries Let p be a natural number. Let - be a p-regular partition of n. The p-rim of - is a part of the rim of - ([JK], p. 56), which is composed of p-segments. Each p-segment except possibly the last contains p points. The first p-segment consists of the first p points of the rim of -, starting with the longest row. (If the rim contains at most p points it is the entire rim.) The next segment is obtained by starting in the row next below the previous p-segment. This process is continued until the final row is reached. We let a 1 be the number of nodes in the p-rim of - (1) and let r 1 be the number of rows in -. Removing the p-rim of - (1) we get a new p-regular partition - (2) of n \Gamma a 1 . We let a be the length of the p-rim and the number of parts of - (2) respectively. Continuing this way we get a sequence of partitions - and a corresponding Mullineux symbol of - a 1 a 2 \Delta \Delta \Delta am The integer m is called the length of the symbol. For p ? n, the well-known Frobenius symbol F (-) of - is obtained from G p (-) as above by As usual, here the top and bottom line give the arm and leg lengths of the principal hooks. It is easy to recover a p-regular partition - from its Mullineux symbol G p (-). Start with the hook - (m) , given by am backwards. In placing each p-rim it is convenient to start from below, at row r i . Moreover, by a slight reformulation of a result in [M1], the entries of G p (-) satisfy (see [AO]) we call the corresponding column \Delta of the Mullineux symbol a singular column, otherwise the column is called regular. If G p (-) is as above then the Mullineux conjugate - M of - is by definition the p-regular partition satisfying a 1 a 2 \Delta \Delta \Delta am In particular, for p ? n, this is just the ordinary conjugation of partitions. Example. Let (In both cases the nodes of the successive 5-rims are numbered 1; 2; 3; 4). Thus Now let p be a prime number and consider the modular representations of S n in characteristic p; note that for all purely combinatorial results the condition of primeness is not needed. The modular irreducible representations D - of S n may be labelled by p-regular partitions - of n, a partition being p-regular if no part is repeated p (or more) times ([JK], 6.1); this is the labelling we will consider in the sequel. Tensoring the modular representation D - of S n by the sign representation of another modular irreducible representation, labelled by a p-regular partition - P . Mullineux has then conjectured [M1]: Conjecture. For any p-regular partition - of n we have - If - is a p-regular partition we let as before a 1 a 2 \Delta \Delta \Delta am denote its Mullineux symbol. We then define the Residue symbol R p (-) of - as oe where x j is the residue of am+1\Gammaj \Gamma r m+1\Gammaj modulo p and y j is the residue of p. Note that the Mullineux symbol G p (-) can be recovered from the Residue symbol R p (-) because of the strong restrictions on the entries in the Mullineux symbol. Also, it is very useful to keep in mind that for a residue symbol there are no restrictions except that (which would correspond to starting with the p-singular partition (1 p )). We also note that a column \Delta in R p (-) is a singular column in G p (-) if and only and R 5 oe Also for the residue symbol of a p-regular partition we have a good description of the residue symbol of its Mullineux conjugate; this is just obtained by translating the definition of the Mullineux map on the Mullineux symbol to the residue symbol notation. Lemma 2.1 Let the residue symbol of the p-regular partition - be oe Then the residue symbol of - M is ae oe where Notation. We now fix a p-regular partition -. Then e - denotes the partition obtained from - by removing all those parts which are equal to 1. We will assume that - has d such parts, 0 - d - - be the partition obtained from - by subtracting 1 from all its parts. We say that - is obtained by removing the first column from -. Unless otherwise specified we assume that the residue symbol R p (-) for - is as above. For later induction arguments we formulate the connection between the residue symbols of - and -. First we consider the process of first column removal; this is an easy consequence of Proposition 1.3 in [BO] and the definition of the residue symbol. Lemma 2.2 Suppose that oe Then ae x 0 oe Here y 0 is defined to be 1 and the - j 's are defined by Moreover, if x then the first column in R p (-) (consisting of x 0 1 and y 0 omitted. Remark 2.3 In the notation of Lemma 2.2 the number d of parts equal to 1 in - is determined by the congruence Moreover since r 1 is the number of parts of - and y is clear that y m is the p-residue of the lowest node in the first column of -. Next we consider the relationship between - and - from the point of adding a column to -; this follows from Proposition 1.6 in [BO]. Lemma 2.4 Suppose that ae x 0 oe Then oe Here x and the - 0 j 's are defined by Moreover, if y 0 then the first column in R p (-) (consisting of Remark 2.5 In the notation of Lemma 2.2 and Lemma 2.4 we have Indeed, (by definition of - j ) (by definition of - 0 3 Mullineux fixed points in a p-block The p-core - (p) of a partition - is obtained by removing p-hooks as long as possible; while the removal process is not unique the resulting p-regular partition is unique as can most easily be seen in the abacus framework introduced by James. The reader is referred to [JK] or [O1] for a more detailed introduction into this notion and its properties. We define the weight w of - by The representation theoretic significance of the p-core is the fact that it determines the p-block to which an ordinary or modular irreducible character labelled by - belongs. The weight of a p-block is the common weight of the partitions labelling the characters in the block. be a partition of n. Then is the Young diagram of -, and (i; called a node of -. If is a node of - and Y (-) n f(i; j)g is again a Young diagram of a partition, then A is called a removable node and - n A denotes the corresponding partition of Similarly, if IN is such that Y (-) [ f(i; j)g is the Young diagram of a partition of n is called an indent node of - and the corresponding partition is denoted - [ A. The p-residue of a node A = (i; j) is defined to be the residue modulo p of denoted res p). The p-residue diagram of - is obtained by writing the p-residue of each node of the Young diagram of - in the corresponding place. The p-content of a partition - is defined by counting the number of nodes of a given residue in the p-residue diagram of -, i.e. c i is the number of nodes of - of p-residue i. In the example above, the p-content of - is It is important to note that the p-content determines the p-core of a partition. This can be explained as follows. First, for given the associated ~n-vector by for any vector there is a unique p-core - with this ~n-vector ~n associated to its p-content c(-) (for short, we also say that ~n is associated to -.) We refer the reader to [GKS] for the description of the explicit bijection giving this relation. From [GKS] we also have the following Proposition 3.1 Let - be a p-core with associated ~n-vector ~n. Then in i with How do we obtain the ~n-vector associated to - from its Mullineux or residue symbol? This is answered by the following Proposition 3.2 Let - be a p-regular partition with Mullineux symbol resp. residue symbol a 1 a 2 \Delta \Delta \Delta am and R p oe Then the associated ~n-vector Proof. In the residue symbol, singular columns do not contribute to the n-vector as they contain the same number of nodes for each residue. So let us consider a regular column y a r \Delta in the Mullineux symbol and the corresponding p-rim in the p-residue diagram. In this case, the contribution only comes from the last section of the p-rim. The final node is in row r and column 1 so its p-residue is What is the p-residue of the top node of this rim section? The length of this section is j a (mod p), hence we have to go j a \Gamma 1 steps from the final node of residue y to the top node of the section, which hence has p-residue j y (going a north or east step always increases the p-residue by 1!). Thus going along the residues in the last section we have a strip the contribution of the intermediate residues to the ~n-vector cancel out, and we only have a contribution 1 for n x and \Gamma1 for n y\Gamma1 , which proves the claim. 2 First we use the preceding proposition to give a short proof of a relation already noticed by Mullineux [M2]: Corollary 3.3 Let - be a p-regular partition. Then Proof. Let the residue symbol of - be R p oe So by Lemma 2.1 we have R p (- M ae oe with Now we consider the contributions of the entries in the residue symbol to the ~n-vectors. If x then we get a contribution 1 to and \Gamma1 to n k (-) on the one hand, and a contribution 1 to n \Gamma(k+1) (- M ) and \Gamma1 to n \Gamma(j+1) (- M ) on the other hand. If x then from column i in the residue symbol we get neither a contribution to ~n(-) nor to ~n(- M ). Hence \Gamman \Gamma(j+1) (-) for all j, i.e. if Now let be the p-content of -, then c(- 0 and hence (\Gamman Now we turn to Mullineux fixed points. Proposition 3.4 Let p be an odd prime and suppose that - is a p-regular partition with - M . Then the representation D - belongs to a p-block of even weight w. Proof. If - M , then its Mullineux symbol is of the form a 1 a 2 \Delta \Delta \Delta am where as before " and where a i is even if and only if pja i . Now by Proposition 3.1 we have is the ~n-vector associated to - and ~ 1). By Proposition 3.2 we have For a we do not get a contribution to the ~n-vector. For a i 6j 0 (mod p) with a i \Gamma1j j (mod p) we get a contribution 1 to n j and \Gamma1 to n \Gamma(j+1) . Note that we can not get any contribution to n p\Gamma1. Thus we have Now we obtain for the weight modulo 2: Hence the weight is even, as claimed. 2 For the following theorem we recall the definition of the numbers k(r; s): is a partition, and r In view of the now proved Mullineux conjecture, the following combinatorial result implies a representation theoretical result in [O2]. Theorem 3.5 Let p be an odd prime. Let - be a symmetric p-core and n 2 IN with even. Then Proof. We set For - 2 F(-) we consider its Mullineux symbol; as - is a Mullineux fixed point this has the form a 1 a 2 \Delta \Delta \Delta am a 1 +" 1a 2 +" being even if and only if pja i . In this special situation the general restrictions on the entries in Mullineux symbols stated at the beginning of x2 are now given by: (ii) If a (iii) If a (iv) a i is even if and only if pja (v) We have already explained before how to read off the p-core of a partition from its Mullineux symbol by calculating the ~n-vector. In the proof of the previous proposition we have already noticed that entries a i j 0 (mod p) do not contribute to the ~n-vector. Now the partitions (a properties (i) to (iv) above are just the partitions satisfying the special congruence and difference conditions for and the congruence set considered in [B], [AO]. The bijection described there transforms the set of partitions above into the set where modN b denotes the smallest positive number congruent to b mod N . Computing the ~n-vector from the b i 's instead of the a i 's with the formula given in the previous proof then gives the same answer since the congruence sequence of the b i 's is the same as the congruence sequence of the regular a i 's. For a bar partition b 2 D as above we then compute its so called N -bar quo- tient; since b has no parts congruent to 0 or p modulo N , the bar quotient is a p\Gamma1-tuple of partitions. For the properties of these objects we refer the reader to [MY], [O1]. It remains to check that the N-weight of b equals w, i.e. that the N -bar core N) of b satisfies We recall from above that we have for the ~n-vector of -: and Hence by Proposition 3.1 we obtain As remarked before the bijection transforming a leaves the congruence sequence modulo of the regular elements in a invariant. Now for determining the N-bar core of b we have to pair off b i 's congruent to 2j each only have to know for each such j the number But this is equal to Now the contribution to the 2p-bar core from the conjugate runners 2j +1 and 2 is for any value of n j easily checked to be Thus the total contribution to the 2p-bar core is exactly the same as the one calculated above, i.e. we have as was to be proved. 2 4 The combinatorial part of the Mullineux conjec- ture We are now going to introduce the main combinatorial concepts for our inves- tigations. The concept of the node signature sequence and the definition of its good nodes have their origin in Kleshchev's definition of good nodes of a partition. First we recapitulate his original definition [K2]. We write the given partition in the form For we then define a t and fl(i; Furthermore, for We then call i normal if and only if for all there exists and such that We call i good if it is normal and if fl(i; Let us translate this into properties of the nodes of - in the Young diagram that can most easily be read off the p-residue diagram of -. One sees immediately that fi(i; j) is just the length of the path from the node at the beginning of the i-th block of - to the node at the end of the jth block of -. The condition is then equivalent to the equality of the p-residue of the indent node in the outer corner of the ith block and the p-residue of the removable node at the inner corner of the jth block. Similarly, is equivalent to the equality of the p-residues of the removable nodes at the end of the ith and jth block. We will say that a node A = (i; j) is above the node below and write this relation as B % A. Then a removable node A of - is normal if for any B 2 indent node of - above A with res res Ag we can choose a removable node CB of - with A % CB % B and res res A, such that jfC j. A node A is good if it is the lowest normal node of its p-residue. Consider the example 5. In the p-residue diagram below we have included the indent node at the beginning of the second block, marked 3, and we have also marked the removable node of residue 3 at the end of the fourth block in boldface. The equality of these residues corresponds to We also see immediately from the diagram below that 4 0The set M i corresponds in this picture to taking the removable node, say A, at the end of the ith block and then collecting into M i (resp. MA ) all the indent nodes above this block of the same p-residue as A. For i resp. A being normal, we then have to check whether for any such indent node, B say, at the end of the jth block we can find a removable node C = CB between A and B of the same p-residue, and such that the collection of all these removable nodes has the same size as M i (resp. MA ). The node A (resp. i) is then good if A is the lowest normal node of its p-residue. The critical condition for the normality of i resp. A above is just a lattice con- dition: it says that in any section above A there are at least as many removable nodes of the p-residue of A as there are indent nodes of the same residue. With these notions the Mullineux conjecture was reduced by Kleshchev to the combinatorial Conjecture. Let - be a p-regular partition, A a good node of -. Then there exists a good node B of the Mullineux image - M such that (-nA) Now we define signature sequences. A (p)-signature is a pair c" where c 2 f0; is a sign. Thus 2+ and 3\Gamma are examples of 5-signatures. A (p)-signature sequence X is a sequence where each c i " i is a signature. Given such a signature sequence X we define for We make the conventions that an empty sum is 0 and that + is counted as +1 in the sum. The i'th peak value - i (X) for X is defined as and the i'th end value defined as We call i a good residue for X if - In that case let and we then say that the residue c k at step k is i-good for X, for short: c k is i-good for X. Let us note that if c k is i-good for X then c Indeed, if is clear since otherwise - contrary to the definition of c k . contrary to the definition of - i (X). The residue c l is called i-normal if c l is i-good for the truncated sequence The following is quite obvious from the definitions. Lemma 4.1 Let X be a signature sequence and let X be obtained from X by adding a signature c t " t at the end. For the following statements are equivalent c t is i-good for X. We are going to define two signature sequences based on -, the node sequence N(-) and the Mullineux sequence M(-). Although they are defined in very different ways we will show that they have the same peak and end value for each i. The node sequence N(-) consists of the residues of the indent and removable nodes of -, read from right to left, top to bottom in -. For each indent residue the sign is + and for each removable residue the sign is \Gamma. Let us note that according to Remark 2.3 the final signature in N(-) is Example. Let we have only indicated the removable and indent nodes in the 5-residue diagram of -.4 Residue Peak value Good (The good signatures (peaks) are underlined and the normal signatures marked with a prime.) In other words, in the node sequence N(-) defined before, if c corresponds to the removable node A, then c and A is normal if and only if the sequence of signs to the left of A belonging to c j 's with c res A is latticed read from right to left. Again, the node A resp. c m is good if it is the last normal node of its residue resp. of its value. The peak value of the node sequence N(-) is the number of normal nodes of -. Remark 4.2 Let as before denote e - the partition obtained from - by removing all those parts which are equal to 1, and let - be the partition obtained from - by subtracting 1 from all its parts. From the definitions it is obvious that for all i Proposition 4.3 Let - and - be as above, and let d be the number of parts 1 in -. and if (3) We have unless the following conditions are all fulfilled In that case y m is i-good for N(-) and Proof. Assume that N(-) consists of m 0 signatures (m 0 is odd). Then N(-) consists of Suppose that If then where in both sequences c m \Gamma. From this and the definition of end values (1) and (2) follows easily. Also since the final sign is \Gamma we have proving (3) in this case. Suppose d 6= 0. If again Again (1) and (2) follows easily. To prove (3) we consider the sequence Obviously for all i. The final signature of N(-) has no influence on - i (N(-)), since the sign is \Gamma. Therefore in order for - i (N (-)) to be different from - i (N(-)) we need 4.1. Thus condition (i) of (3) is fulfilled and condition (iii) follows from ( ). Since by assumption d 6= 0 (ii) is also fulfilled. Thus (3) is proved in this case also. 2 We proceed to prove an analogue of Proposition 4.3 for the Mullineux (signa- ture) sequence M(-), which is defined as follows: Let the residue symbol of - be oe Then Starting with the signature 0\Gamma corresponds to starting with an empty partition at the beginning which just has the indent node (1; 1) of residue 0. oe and (The good signatures in M(-) are again underlined and the normal signatures marked with a prime.) The table above is identical with the one in the previous example. Lemma 4.4 Let - and - be as above. Let M (-) be the signature sequence obtained from M(-) by removing the two final signatures y m+ and (y Then for all i we have Proof. We use the notation of Lemma 2.2 for R p (-) and R p (-) and proceed by induction on m. First we study the beginnings of M (-) and M(-). We compare (from M (-)) with We have put brackets [ ] around a part of (2), because these signatures do not occur when x The former gives a contribution \Gamma1 to residue 1 and contributions 0 to all others, the latter a contribution \Gamma1 to residue 0 and 0 to all others. The signatures 0\Gamma 0+ in the latter sequence have no influence on the end values and peak values of M(-), (even when x may be ignored. Then again we see that (1) gives the same contribution to residue i as (2) 0 to residue our result is true. We assume that the result is true for partitions whose Mullineux symbols have length and have to compare (from M (-)) with (from M(-)) By Lemma 2.2, (4) may be written as We see that up to rearrangement the difference between the residues occurring in (3) and (4) 0 is just ffi m . Whereas the rearrangement is irrelevant for the end values it could influence the peak value if signatures with same residue but different signs are interchanged. The possible coincidences of residues with different signs are (first and fourth residue in (3)) or (second and third residue in (3)) But the equations (ff) and (fi) are equivalent and by Lemma 2.2 they are fulfilled if and only if becomes and In this case the difference between the occurring residues in 1 (without rear- rangement) and our statement is true. If y then the difference between the occurring residues is again 1(= since there is no coincidence for residues with different signs we may apply Lemma 4.1 and the induction hypotheses to prove the statement in this case too. 2 Lemma 4.5 Suppose that in the notation as above we have for Then d 6= 0. Proof. Suppose ends on (i \Gamma 1)\Gamma. But then clearly ! Lemma 4.6 Let the notation be as in Lemma 4.4. (1) For y j is i-good for M (-) ae j+1 is (i \Gamma 1)-good for M(-) or j+1 is (i \Gamma 1)-good for M(-) (2) For x j is i-good for M (-) j is (i \Gamma 1)-good for M(-) Proof. This follows immediately from the proof of Lemma 4.4. It should be noted that x 1 cannot be 0-good for M (-) since M (-) starts by 0\Gamma. More- over, the proof of Lemma 4.4 shows that if x j is i-good for M (-), then we cannot have Proposition 4.7 Let - and - be as above. and if (3) We have In that case y m is i-good for M(-) and Note. There is a strict analogy between the Propositions 4.3 and 4.7. In part (3) the assumption d 6= 0 is not necessary in Proposition 4.7 due to Lemma 4.5. Proof. By Lemma 4.4 If we add y m+ and (y m \Gamma 1)\Gamma to M (-) we get M(-). Therefore an argument completely analogous to the one used in the case d 6= 0 in the proof of Proposition 4.3 may be applied. 2 Theorem 4.8 Let - be a p-regular partition. Then for all Proof. We use induction on the number ' of columns in -. For d and R p ae 0 oe . Thus and the result is clear. Assume the result has been proved for partitions with 2. Let - be obtained by removing the first column from -. By the induction hypothesis we have for all i. Using Proposition 4.3 and Proposition 4.7 (see also the note to Proposition 4.7) we get the result. 2 Theorem 4.9 The following statements are equivalent for a p-regular partition (1) There is a good node of residue i in -. (2) M(-) has i as a good residue. (3) N(-) has i as a good residue. Proof. (1) , (3): See the beginning of this section. (2) , (3): Theorem 4.8. 2 Finally we describe the effect of the removal of a good node on the residue (or equivalently on the Mullineux symbol). First we prove a lemma: Lemma 4.10 Suppose that there is a good node of residue i in -. Then the following statements are equivalent: (1) The good node of residue i occurs in the first column of -. Proof. The statement (1) clearly is equivalent to (where as before e - is obtained from - by removing all parts equal to 1.) We now have (by Theorem 4.8) is i-good for M(-) (by Proposition 4.7)Theorem 4.11 Suppose that the p-regular partition - has a good node A of residue i. Let oe Then for some j, 1 - j - m, one of the following occurs: (1) x j is i-good for M(-) and oe (2) y j is i-good for M(-) and oe resp. oe Proof. The proof is by induction on j-j. Suppose first that A occurs in the first column of -. Then the first column in G p (-nA) is obtained from the first column in G p (-) by subtracting 1 in each entry and all other entries are unchanged; note that in the case where G p , we have a degenerate case and G p (-nA) is the empty symbol. By definition of the residue symbol this means that y m in R p (-) is replaced by y of course, in the degenerate case also R p (-nA) is the empty residue symbol. On the other hand y m is i-good for M(-) by Lemma 4.10, and in the degenerate case y 1 is 0-good for M(-), and so we are done in this case. Now we assume that A does not occur in the first column of -. Let B be the node of - corresponding to A. Clearly B is a good node of residue We may apply the induction hypothesis to - and B. Suppose that ae x 0 oe By the induction hypothesis we know that one of the following cases occurs: Case I. x 0 j in R p (-) is replaced by x 0 j is (i \Gamma 1)-good for M(-). Case II. y 0 j in R p (-) is replaced by y 0 j is (i \Gamma 1)-good for M(-), resp. in the degenerate case y 0 1 is 0-good for M(-) and then the first column x 0y 0in R p (-) is omitted in R p (-nB). We treat Case I in detail. Case II is treated in a similar way. Case I: By Lemma 4.6 we have one of the following cases: We add a first column to -nB to get -nA. Then R p (-nA) is obtained from using Lemma 2.4. We fix the notation ae x 00 oe and oe Case Ia. We know x 0 since we are in Case I and y since we are in Case Ia. Moreover since Also x 0 2.2. By Lemma 2.4 where ae 0 if x 00 But x 00 by the above. Thus 1. It is readily seen that all other entries in R p (-nA) coincide with those of R p (-). Thus possibility (2) occurs in the theorem. Case Ib. We know x 0 since we are in Case I and x since we are in Case Ib. Also since i.e. x 00 2. Let ffi 00 again be defined by ae 0 if x 00 Then by Lemma 2.4 x . We claim that ffi 00 j and we also know that x 0 by definition of M(-) x 0 j is not a peak, contrary to our assumption that we are in Case I. Thus ffi 00 Again it is easily seen that all other entries of R p (-nA) coincide with those of R p (-). Thus possibility (1) occurs in the theorem. 2 We illustrate the theorem above by giving Kleshchev's p-good branching graph for p-regular partitions for in both the usual and the residue symbol notation; we recall that the p-good branching graph for p- regular partitions is also the crystal graph for the basic representation of the quantum affine algebra (see [LLT] for these connections). Below, an edge from a partition - of m to a partition - of labelled by the residue r if - is obtained from - by removing a node of residue those edges are drawn that correspond to the removal of good nodes. @ @ @ @ @ @ @ @ @ @ @ @ @ @ @ @ @ @ @ @ @ @ @ @ We can now easily deduce the combinatorial conjecture to which the Mullineux conjecture had been reduced by Kleshchev: Corollary 4.12 Suppose that the p-regular partition - has a good node A of residue i. Then its Mullineux conjugate - M has a good node B of residue \Gammai satisfying Proof. Considering the residue symbol of - it is easily seen that the Mullineux sequence of - and its conjugate - M are very closely related. Indeed, the peak and end values for each residue i in M(-) equal the corresponding values for the residue \Gammai in M(- M ), and if there is an i-good node at column k in the residue symbol of -, then there is an \Gammai-good node at column k in the residue symbol of - M . More precisely, in the regular case these good nodes are one at the top and one at the bottom of the column, whereas in the singular case both are at the top. Comparing this with Theorem 4.11 implies the result. 2 Acknowledgements . The authors gratefully acknowledge the support of the Danish Natural Science Foundation and of the EC via the European Network 'Algebraic Combinatorics' (grant ERBCHRXCT930400). --R Partition identities with an application to group representation theory A combinatorial proof of a refinement of the Andrews- Olsson partition identity Theory (A) 68 A proof of the Mullineux conjecture Cranks and t-cores The representation theory of the symmetric group Branching rules for modular representations of symmetric groups I Branching rules for modular representations of symmetric groups II Branching rules for modular representations of symmetric groups III On decomposition numbers and branching coefficients for symmetric and special linear groups Hecke algebras at roots of unity and crystal basis of quantum affine algebras Crystal base of the basic representation of U q Some combinatorial results involving Young diagrams Bijections of p-regular partitions and p-modular irreducibles of symmetric groups On the p-cores of p-regular diagrams Combinatorics and representations of finite groups The number of modular characters in certain blocks --TR --CTR Jonathan Brundan , Jonathan Kujawa, A New Proof of the Mullineux Conjecture, Journal of Algebraic Combinatorics: An International Journal, v.18 n.1, p.13-39, July

signature sequence;good nodes in residue diagram;modular representation;mullineux conjecture;symmetric group

291689

The Enumeration of Fully Commutative Elements of Coxeter Groups.

A Coxeter group element w is fully commutative if any reduced expression for w can be obtained from any other via the interchange of commuting generators. For example, in the symmetric group of degree n, the number of fully commutative elements is the nth Catalan number. The Coxeter groups with finitely many fully commutative elements can be arranged into seven infinite families A_n, B_n, D_n, E_n, F_n, H_n and I_2(m). For each family, we provide explicit generating functions for the number of fully commutative elements and the number of fully commutative involutions; in each case, the generating function is algebraic.

Introduction A Coxeter group element w is said to be fully commutative if any reduced word for w can be obtained from any other via the interchange of commuting generators. (More explicit definitions will be given in Section 1 below.) For example, in the symmetric group of degree n, the fully commutative elements are the permutations with no decreasing subsequence of length 3, and they index a basis for the Temperley-Lieb algebra. The number of such permutations is the nth Catalan number. In [St1], we classified the irreducible Coxeter groups with finitely many fully commutative elements. The result is seven infinite families of such groups; namely, An , Bn , Dn , En , Fn , Hn and I 2 (m). An equivalent classification was obtained independently by Graham [G], and in the simply-laced case by Fan [F1]. In this paper, we consider the problem of enumerating the fully commutative elements of these groups. The main result (Theorem 2.6) is that for six of the seven infinite families (we omit the trivial dihedral family I 2 (m)), the generating function for the number of fully commutative elements can be expressed in terms of three simpler generating functions for certain formal languages over an alphabet with at most four letters. The languages in question vary from family to family, but have a uniform description. The resulting generating function one obtains for each family is algebraic, although in some cases quite complicated. (See (3.7) and (3.11).) In a general Coxeter group, the fully commutative elements index a basis for a natural quotient of the corresponding Iwahori-Hecke algebra [G]. (See also [F1] for the simply- laced case.) For An , this quotient is the Temperley-Lieb algebra. Recently, Fan [F2] has shown that for types A, B, D, E and (in a sketched proof) F , this quotient is generically semisimple, and gives recurrences for the dimensions of the irreducible representations. the question of semisimplicity remains open.) This provides another possible approach to computing the number of fully commutative elements in these cases; namely, as the sum of the squares of the dimensions of these representations. Interestingly, Fan also shows that the sum of these dimensions is the number of fully commutative involutions. With the above motivation in mind, in Section 4 we consider the problem of enumerating fully commutative involutions. In this case, we show (Theorem 4.3) that for the six nontrivial families, the generating function can be expressed in terms of the generating functions for the palindromic members of the formal languages that occur in the unrestricted case. Again, each generating function is algebraic, and in some cases, the explicit form is quite complicated. (See (4.8) and (4.10).) In Section 5, we provide asymptotic formulas for both the number of fully commutative elements and the number of fully commutative involutions in each family. In an appendix, we provide tables of these numbers up through rank 12. 1. Full Commutativity Throughout this paper, W shall denote a Coxeter group with (finite) generating set S, Coxeter graph \Gamma, and Coxeter matrix standard reference is [H]. 1.1 Words. For any alphabet A, we use A to denote the free monoid consisting of all finite-length words A. The multiplication in A is concatenation, and on occasion will be denoted ' \Delta '. Thus (a; b)(b; a) subsequence of a obtained by selecting terms from a set of consecutive positions is said to be a subword or factor of a. For each w 2 W , we define R(w) ae S to be the set of reduced expressions for w; i.e., the set of minimum-length words such that For each integer m - 0 and s; t 2 S, we define and let - denote the congruence on S generated by the so-called braid relations; namely, for all s; t 2 S such that m(s; t) ! 1. It is well-known that for each w 2 W , R(w) consists of a single equivalence class relative to -. That is, any reduced word for w can be obtained from any other by means of a sequence of braid relations [B,xIV.1.5]. 1.2 Commutativity classes. Let - denote the congruence on S generated by the interchange of commuting gener- ators; i.e., (s; t) - (t; s) for all s; t 2 S such that m(s; 2. The equivalence classes of this congruence will be referred to as commutativity classes. Given the heap of s is the partial order of f1; obtained from the transitive closure of the relations i OE j for all 3. It is easy to see that the isomorphism class of the heap is an invariant of the commutativity class of s. In fact, although it is not needed here, it can be shown that only if there is an isomorphism i 7! i 0 of the corresponding heap orderings with s example, see Proposition 1.2 of [St1].) In [St1], we defined w 2 W to be fully commutative if R(w) consists of a single commutativity class; i.e., any reduced word for w can be obtained from any other solely by use of the braid relations that correspond to commuting generators. It is not hard to show that . I (m)54 F A Figure 1: The irreducible FC-finite Coxeter groups. this is equivalent to the property that for all s; t 2 S such that m(s; t) - 3, no member of R(w) has hs; ti m as a subword, where It will be convenient to let W FC denote the set of fully commutative members of W . As mentioned in the introduction, the irreducible FC-finite Coxeter groups (i.e., Coxeter groups with finitely many fully commutative elements) occur in seven infinite families denoted An , Bn , Dn , En , Fn , Hn and I 2 (m). The Coxeter graphs of these groups are displayed in Figure 1. It is interesting to note that there are no "exceptional" groups. For the dihedral groups, the situation is quite simple. Only the longest element of I 2 (m) fails to be fully commutative, leaving a total of 2m \Gamma 1 such elements. Henceforth, we will be concerned only with the groups in the remaining six families. 1.3 Restriction. For any word s 2 S and any J ae S, let us define sj J to be the restriction of s to J ; i.e., the subsequence formed by the terms of s that belong to J . Since the interchange of adjacent commuting generators in s has either the same effect or no effect in sj J , it follows that for any commutativity class C, the restriction of C to J is well-defined. Figure 2: A simple branch. A family F of subsets of S is complete if for all s 2 S there exists J 2 F such that and for all s; t 2 S such that m(s; t) - 3 there exists J 2 F such that s; t 2 J . Proposition 1.1. If F is a complete family of subsets of S, then for all s; s 0 2 S , we have Proof. The necessity of the stated conditions is clear. For sufficiency, suppose that s is the first term of s. Since s 2 J for some J 2 F , there must also be at least one occurrence of s in s 0 . We claim that any term t that precedes the first s in s 0 must commute with s. If not, then we would have sj fs;tg 6- s 0 j fs;tg , contradicting the fact that sj J - s 0 j J for some J containing fs; tg. Thus we can replace s 0 with some s 00 - s 0 whose first term is s. If we delete the initial s from s and s 00 , we obtain words that satisfy the same restriction conditions as s and s 0 . Hence s - s 00 follows by induction with respect to length. 2. The Generic Case Choose a distinguished generator s 1 2 S, and let denote the infinite sequence of Coxeter groups in which W i is obtained from W i\Gamma1 by adding a new generator s i such that m(s commutes with all other generators of W i\Gamma1 . In the language of [St1], fs is said to form a simple branch in the graph of Wn . For ng denote the generating set for Wn , and let \Gamma n denote the corresponding Coxeter graph. (See Figure 2.) It will be convenient also to let S 0 and denote the corresponding data for the Coxeter group W 0 obtained when s 1 is deleted from S. Thus 2.1 Spines, branches, and centers. For any w 2 W FC n , we define the spine of w, denoted oe(w), to be the pair (l; A), where l denotes the number of occurrences of s 1 in some (equivalently, every) reduced word for w, and A is the subset of defined by the property that k 2 A iff there is no occurrence of s 2 between the kth and (k 1)th occurrences of s 1 in some (equivalently, reduced word for w. We refer to l as the length of the spine. Continuing the hypothesis that w is fully commutative, for J ' Sn we let wj J denote the commutativity class of sj J for any reduced word s 2 R(w). (It follows from the Figure 3: An F 7 -heap. Figure 4: Center and branch. discussion in Section 1.3 that this commutativity class is well-defined.) In particular, for each w 2 W FC n , we associate the pair (wj Sn \GammaS We refer to wj Sn \GammaS 0 and wj S1 as the branch and central portions of w, respectively. For example, consider the Coxeter group F 7 . We label its generators fu; t; s in the order they appear in Figure 1, so that fs is a simple branch. The heap of a typical fully commutative member of F 7 is displayed in Figure 3. Its spine is (5; f1; 4g), and the heaps of its central and branch portions are displayed in Figure 4. "branch set") to be the set of all commutativity classes B over the ng such that a subword of some member of B, then a subword of some member of B, then Furthermore, given a spine oe = (l; A), we define Bn (oe) to be the set of commutativity Bn such that there are l occurrences of s 1 in every member of B, and (B3) The kth and (k 1)th occurrences of s 1 occur consecutively in some member of B if and only if k 2 A. We claim (see the lemma below) that Bn (oe) contains the branch portions of every fully commutative w 2 Wn with spine oe. Note also that Bn depends only on n, not W . Similarly, let us define "central set") to be the set of commutativity classes C over the alphabet S (C1) For all s 2 S 1 , no member of C has (s; s) as a subword. (C2) If hs; ti m is a subword of some member of C, where occurs at least twice in this subword. (In particular, s In addition, we say that compatible with the spine every member of C has l occurrences of s 1 , and is a subword of some member of C, where subword includes the kth and (k 1)th occurrences of s 1 for some k 62 A. denote the set of oe-compatible members of C. We claim (again, see the lemma below) that C(oe) contains the central portions of every w 2 W FC n with spine oe. Note also that C(oe) depends only on (more precisely, on the Coxeter graph \Gamma), not the length of the branch attached to it. Lemma 2.1. The mapping w 7! (wj Sn \GammaS defines a bijection oe Bn (oe) \Theta CW (oe): Proof. For all non-commuting pairs s; t 2 Sn , we have fs; tg so by Proposition 1.1, the commutativity class of any w 2 W FC (and hence w itself) is uniquely determined by wj Sn \GammaS 0 and wj S1 . Thus the map is injective. Now choose an arbitrary fully commutative w 2 Wn with spine . The defining properties of the spine immediately imply the validity of (B3). Since consecutive occurrences of any s 2 Sn do not arise in any s 2 R(w), it follows that for all k - 1, the kth and (k 1)th occurrences of s in s must be separated by some t 2 Sn such that m(s; t) - 3. For the only possibilities for t are in Sn holds. For s 2 S 0 , the only possibilities for t are in S 1 , so (C1) could fail only if some k, the only elements separating the kth and (k +1)th occurrences of s 1 in s that do not commute with s 1 are one or more occurrences of s 2 . In that case, we could choose a reduced word for w so that the subword running from the kth to the (k +1)th occurrences of s 1 forms a reduced word for a fully commutative element of the parabolic subgroup isomorphic to An generated by fs g. However, it is easy to show (e.g., Lemma 4.2 of [St1]) that every reduced word for a fully commutative member of An has at most one occurrence of each "end node" generator. Thus (C1) holds. Concerning (B2), (C2) and (C3), suppose that (s occurs as a subword of some member of the commutativity class B. If i ? 1, then every s 2 Sn that does not commute with s i belongs to Sn \Gamma S 0 . Hence, some reduced word for w must also contain the subword contradicting the fact that w is fully commutative. Thus (B2) holds. Similarly, if we suppose that hs; ti m occurs as a subword of some member of C, where and s; t 2 S 1 , then again we contradict the hypothesis that w is fully commutative unless is the only member of S 1 that may not commute with some member of Sn In either case, since hs; ti m cannot be a subword of any s 2 R(w), it follows that s 1 occurs at least twice in hs; ti m (proving (C2)), and between two such occurrences of s 1 , say the kth and (k + 1)th, there must be an occurrence of s 2 in s. By definition, this means k 62 A, so (C3) holds. Thus B 2 Bn (oe) and C 2 CW (oe). Finally, it remains to be shown that the map is surjective. For this, let be a spine, and choose commutativity classes B 2 Bn (oe) and C 2 CW (oe). Select representatives is a singleton, and this singleton appears the same number of times in s B and s C (namely, l times), it follows that there is a word s 2 S whose restrictions to Sn \Gamma S 0 and S 1 are s B and s C , respectively. We claim that s is a reduced word for some w 2 W FC n , and hence w 7! (B; C). To prove the claim, first consider the possibility that for some s 2 Sn , (s; s) occurs as a subword of some member of the commutativity class of s. In that case, depending on whether s 2 S 1 , the same would be true of either B or C, contradicting (B1) or (C1). Next consider the possibility that hs; ti m occurs as a subword of some word s 0 in the commutativity class of s, where 3. We must have either s; t 2 Sn \Gamma S 0 or hence the same subword appears in some member of B or C, respectively. In the former case, (B2) requires that 3. However the restriction of s 0 to S 1 would then have consecutive occurrences of s 1 , contradicting (C1). In the latter case, (C2) and (C3) require that s t, and that the subword hs; ti m includes the kth and (k 1)th occurrences of s 1 in s 0 for some k 62 A. It follows that s 2 does not occur between these two instances of s 1 in s 0 , and thus they appear consecutively in the restriction of s 0 to Sn \Gamma S 0 , contradicting (B3). Hence the claim follows. The above lemma splits the enumeration of the fully commutative parts of the Coxeter groups into two subproblems. The first subproblem, which is universal for all Coxeter groups, is to determine the number of branch commutativity classes with spine oe; i.e., the cardinality of Bn (oe) for all integers n - 0 and all oe. The second subprob- lem, which needs only to be done once for each series Wn , is to determine the number of central commutativity classes with spine oe; i.e., the cardinality of CW (oe). s A ss s Figure 5. 2.2 Spinal analysis. The possible spines that arise in the FC-finite Coxeter groups are severely limited. To make this claim more precise, suppose that is one of the six nontrivial families of FC-finite Coxeter groups (i.e., A, B, D, E, F , or H). The Coxeter graph of W can then be chosen from one of the six in Figure 5. For convenience, we have used s as the label for the distinguished generator previously denoted s 1 . Lemma 2.2. If C 2 CW is compatible with the spine is one of the Coxeter groups in Figure 5, then A ' f1; l \Gamma 1g. Proof. Let s 2 S be a representative of C, and towards a contradiction, let us suppose that A includes some k such that 1. Note that it follows that the kth and 1)th occurrences of s in s are neither the first nor the last such occurrences. For the H-graph, property (C1) implies that the occurrences of s and t alternate in s. Hence, the kth and (k 1)th occurrences of s appear in the middle of a subword of the form (s; t; s; t; s; t; s). In particular, these two occurrences of s participate in a subword of s of the form (t; s; t; s; contradicting (C3). For the F -graph, property (C1) implies that any two occurrences of s must be separated by at least one t. On the other hand, the subword between two occurrences of s must be a reduced word for some fully commutative member of the subgroup generated by ft; ug (property (C2)), so the occurrences of s and t must alternate, and in the restriction of s to fs; tg, the kth and (k 1)th occurrences of s appear in the middle of a subword of the form (s; t; s; t; s; t; s). By (C3), these two occurrences of s cannot participate in an occurrence of (t; s; t; s) or (s; t; s; t) in s. Hence, the two occurrences of t surrounding the occurrence of s must be separated by an occurrence of u. However in that case, (u; t; u) is a subword of some member of the commutativity class of s, contradicting (C2). For the E-graph, at least one of t and t 0 must appear between any two occurrences of s (otherwise (C1) is violated), and both t and t 0 must appear between the kth and 1)th occurrences of s, by (C2). On the other hand, property (C3) also implies that the subword (strictly) between the 2)th occurrences of s in s must be a reduced word for some fully commutative member of W , a Coxeter group isomorphic to A 4 . In particular, this implies that t 0 can appear at most once, and t at most twice, in this subword. Since we have already accounted for at least four occurrences of t and t 0 , we have a contradiction. This completes the proof, since the remaining three graphs are subgraphs of the preceding ones. 2.3 Branch enumeration. The previous lemma shows that for the FC-finite Coxeter groups, we need to solve the branch enumeration problem (i.e., determine the cardinality of Bn (oe)) only for the spines 1g. For this, we first introduce the notation for the number of (n+ sequences. That is, B n;l is the number of orderings of votes for two candidates so that the winning candidate never trails the losing candidate, with the final tally being n l votes to n \Gamma l votes. (For example, see [C, x1.8].) This quantity is also the number of standard Young tableaux of shape (n Lemma 2.3. For integers n; l - 0, we have Proof. For n;l denote the cardinality of Bn (oe) for respectively. In the case oe = (l; ?), the defining properties (B1) and (B3) for membership of B in Bn (oe) can be replaced with member of B has (s It follows that for 1 l, the kth and (k 1)th occurrence of s 1 in any member of B must be separated by exactly one s 2 , and the total number of occurrences of s 2 must be according to whether the first and last occurrences of s 1 are preceded (resp., followed) by an s 2 . Furthermore, the restriction of B to fs ng is a commutativity class with no subwords of the form (s possibly shifting indices (i i), we thus obtain any one of the members of l 0 denotes the number of occurrences of s 2 . Accounting for the four possible ways that s 1 and s 2 can be interlaced (or two, if l = 0), we obtain the recurrence n\Gamma1;l +B (0) On the other hand, it is easy to show that B n;l satisfies the same recurrence and initial . (In fact, one can obtain a bijection with ballot sequences by noting that the terms of the recurrence correspond to specifying the last two votes.) By word reversal, the cases corresponding to are clearly equivalent, so we restrict our attention to the former. Properties (B1) and (B3) imply that the restriction of any B in Bn (oe) to fs must then take the form where each ' ' represents an optional occurrence of s 2 . We declare the left side of B to be open if the above restriction has the form (s there is no s 3 separating the first two occurrences of s 2 . Otherwise, the left side is closed. Case I. The left side is open. In this case, if we restrict B to fs (and shift indices), we obtain any one of the members of according to whether there is an occurrence of s 2 following the last s 1 . (If l = 2, then there is no choice: l is the only possibility.) Case II. The left side is closed. In this case, if we delete the first occurrence of s 1 from B, we obtain any one of the commutativity classes in Bn (l \Gamma The above analysis yields the recurrence 2: It is easy to verify that the claimed formula for B (1) n;l satisfies the same recurrence and the proper initial conditions. For the restriction of any B in Bn (oe) to fs takes the form where again each ' ' represents an optional occurrence of s 2 . In the special case l = 3, this becomes deleting one of the occurrences of s 1 , we obtain any one of the commutativity classes in Bn (2; f1g). Assuming l - 4, we now have not only the possibility that the left side of B is open (as in the case but the right side may be open as well, mutatis mutandis. Case I. The left and right sides of B are both open. In this case, if we restrict B to ng (and shift indices), we obtain any one of the members of Case II. Exactly one of the left or right sides of B is open. Assuming it is the left side that is open, if we restrict B to fs ng (and shift indices), we obtain any one of the members of according to whether there is an occurrence of s 2 following the last s 1 . Case III. The left and right sides of B are both closed. In this case, if we delete the first and last s 1 from B, we obtain any one of the members of Bn (l \Gamma 2; ?). The above analysis yields B (2) n;2 and the recurrence for Once again, it is routine to verify that the claimed formula for B (2) n;l satisfies the same recurrence and initial conditions. Remark 2.4. The union of Bn (l; ?) for all l - 0 is the set of commutativity classes corresponding to the fully commutative members of the Coxeter group Bn whose reduced words do not contain the subword In the language of [St2], these are the "fully commutative top elements" of Bn ; in the language of [F1], these are the "commutative elements" of the Weyl group Cn . Let R(x) denote the generating series for the Catalan numbers. That is, Note that xR(x) 1. The following is a standard application of the Lagrange of [GJ]). We include below a combinatorial proof. Lemma 2.5. We have Proof. A ballot sequence in which A defeats B by 2l votes can be factored uniquely by cutting the sequence after the last moment when candidate B trails by i votes, 1. The first part consists of a ballot sequence for a tie vote, and all remaining parts begin with a vote for A, followed by a ballot sequence for a tie. After deleting the 2l votes for A at the beginnings of these parts, we obtain an ordered 1)-tuple of ballot sequences for ties, for which the generating series is R(x) 2l+1 . 2.4 The generic generating function. To enumerate the fully commutative elements of the family remains is the "central" enumeration problem; i.e., determining the cardinalities of CW (oe) for all spines oe of the form described in Lemma 2.2. Setting aside the details of this problem until Section 3, let us define and let C i denote the generating series defined by l-0 l-3 l-3 Although these quantities depend on W , we prefer to leave this dependence implicit. Theorem 2.6. If W is one of the six Coxeter groups displayed in Figure 5, we have Proof. Successive applications of Lemmas 2.1, 2.2, and 2.3 yield oe l-0 l-3 l-3 n;l l-0 l-3 c l;1 l-3 c l;2 Using Lemma 2.5 to simplify the corresponding generating function, 1 we obtain l-0 l-3 c l;1 l-3 c l;2 It should be noted that when \Gamma1, the coefficient of c l;2 in (2.1) is zero. Thus the range of summation for this portion of the generating function can be extended to n - \Gamma1. Bearing in mind that R(x) is routine to verify that this agrees with the claimed expression. Remark 2.7. As we shall see in the next section, for each series Wn the generating functions C i (x) are rational, so the above result implies that the generating series for belongs to the algebraic function field 3. Enumerating the Central Parts In this section, we determine the cardinalities of the central sets for each of the six Coxeter groups W displayed in Figure 5. (The reader may wish to review the labeling of the generators in these cases, and recall that the distinguished generator s 1 has been given the alias s.) We subsequently apply Theorem 2.6, obtaining the generating function for the number of fully commutative elements in Wn . 3.1 The A-series. In this case, s is a singleton generator, so there is only one commutativity class of each length. It follows easily from the defining properties that the only central commutativity classes are those of (s) and ( ) (the empty word). These are compatible only with the spines respectively. Thus we have and Theorem 2.6 implies Extracting the coefficient of x n , we obtain a result first proved in [BJS, x2]. 3.2 The B-series. In this case, we have and the defining properties imply that the central commutativity classes are singletons in which the occurrences of s and t alternate. It follows that c l;0 is simply the number of alternating fs; tg-words in which s occurs l times; namely, 4 (if l ? Also, the only alternating fs; tg-word that is compatible with a spine (l; A) with A 6= ? is (s; t; s), which is compatible with (2; f1g). Thus we have 4x After some simplifications, Theorem 2.6 yields Extracting the coefficient of x a result first proved in [St2,x5]. 3.3 The D-series. In this case, a set of representatives for the central commutativity classes consist of the subwords of (s; t; s; t Of these, only (s; t; t compatible with a spine (l; A) with A 6= ?; the remainder are compatible only with (l; ?) for some l. Among the subwords of (s; t; s; t number with l occurrences of s is 8 (if l - 2), 7 (if l = 1), or 3 (if l = 0). Thus we have and after some simplifications, Theorem 2.6 implies Extracting the coefficient of x n\Gamma2 , we obtain a result obtained previously in [F1] and [St2,x10]. 3.4 The H-series. As in the B-series, the central commutativity classes are the singletons formed by each of the alternating fs; tg-words. In particular, the value of C 0 (x) is identical to its B-series version. The words that are compatible with spines of the form (l; f1g) are those that begin with s (and have at least two occurrences of s), and (t; s; t; s); thus c 3. The words compatible with spines of the form f1; l \Gamma 1g are those that both begin and end with s and have at least four occurrences of s; i.e., c 4. Thus we have After some simplifications, Theorem 2.6 yields Extracting the coefficient of x 3: (3.4) 3.5 The F -series. In this case, we can select a canonical representative s 2 S from each central commutativity class by insisting that whenever s and u are adjacent in s, u precedes s. Any such word has a unique factorization l each being words consisting of an initial s followed by a ft; ug-word. In fact, given our conventions, we must have with allowed only if We also cannot have (s; t; u) preceded by (u), (t; u), or (s; t; u); otherwise, some member of the commutativity class of s contains the forbidden subword (u; t; u). Conversely, any word meeting these specifications is the canonical representative of some central commutativity class. The language formed by these words therefore consists of together with the exceptional cases f(u); (t; u); (u; s); (t; u; s)g. Hence Turning now to C 1 (x), note that the central commutativity classes that are compatible with a spine of the form (l; f1g) are those for which the first two occurrences of s do not participate in an occurrence of the subwords (s; t; s; t), or (t; s; t; s). If s occurs three or more times, this requires ( ) to be the first factor in (3.5), followed by an occurrence of (s; t; u; s; t). Hence, the canonical representatives compatible with (l; f1g) consist of and four additional cases with l = 2: f(s; t; s); (u; s; t; s); (s; t; s; u); (t; u; s; t; s)g. It follows that c l-3 c l;1 x To determine C 2 (x), note first that (s; t; u; s; t; s) is the unique canonical representative compatible with the spine (3; f1; 2g). For the spines (l; f1; l \Gamma 1g) with l - 4, compatibility requires (s) to be the last factor in (3.6), and it must be preceded by (s; t; u; s; t). Hence l-3 c l;2 x After simplifying the generating function provided by Theorem 2.6, we obtain While it is unlikely that there is a simple closed formula for is interesting to note that the Fibonacci numbers f n satisfy f 2n x f 3n x so when the coefficient of x n\Gamma2 is extracted in (3.7), we obtain 3.6 The E-series. We claim that there is a unique member of each central commutativity class (in fact, any commutativity class in S ) with the property that (s; u), do not occur as subwords. To see this, note first that the set of left members of these pairs is disjoint from the set of right members. Secondly, these pairs are precisely the set of commuting generators of W . Hence, for any pair of words that differ by the interchange of two adjacent commuting generators, one member of the pair can be viewed as a "reduction" of the other, in the sense that the set of positions where u and t occur are farther to the left. Furthermore, since the set of instances of the forbidden pairs in any given word are pairwise disjoint, it follows by induction that any sequence of reductions eventually terminates with the same word, proving the claim. Let L denote the formal language over the alphabet S formed by the canonical representatives (in the sense defined above) of the central commutativity classes. Given any formal language K over S, we will write K(x) for the generating function obtained by assigning the weight x l to each s 2 K for which s occurs l times. Note that by this convention, we have Any word s 2 S has a unique factorization l each being words consisting of an initial s followed by a ft; t 0 ; ug-word. For membership in L, every subword of s not containing s must be a member of the set of canonical representative for the fully commutative members of the subgroup generated by ft; t 0 ; ug. When s is prepended to these words, only six remain canonical: Thus we have For each e 2 E, let L e denote the set of s 2 L for which the initial factor s 0 is e. If or deletion of the initial s in s yields a member of L e for some e 2 f(t); (t; conversely. In terms of generating functions, we have Similarly, deletion of s from the second position defines a bijection from L (u) \Gamma f(u); (u; s)g to so we have Combining these two decompositions, we obtain Now consider the language and the refinements K i consisting of those nonvoid members of K whose initial factor is a i . Since the result of appending a to any s 2 L remains in L if and only if s does not already end in s, it follows that L ( g. Similarly, we have so (3.8) can be rewritten in the form For the commutativity classes of a 2 a 3 ; a 3 a 4 a 3 ; a i a each have representatives in which one or more of the subwords (t; s; and (t; u; t) appear, and hence cannot be central. Conversely, as a subset of fa membership in K is characterized by avoidance of the subwords listed above. It follows that K Solving this recursive description of the languages K i (essentially a computation in the ring of formal power series in noncommuting variables a a 4 a 3 ; a 2 a 4 a 3 a 4 g \Delta f( ); a 2 ; a 2 a 4 a 3 a 5 g. Thus and hence (3.9) implies The central commutativity classes compatible with spines of the form (l; f1g) are those for which the first two occurrences of s do not participate in an occurrence of the subwords These correspond to the members of L for which the first occurrence of one of the factors a i is either a 5 or a 6 , followed by at least one more occurrence of . If a 5 is the first factor, the possibilities are limited to f( ); (u); (t; u)ga 5 a 1 , since a 5 can be followed only by a 1 . If the first factor is a 6 , then the choices consist of the members of K 6 f( ); a 1 g other than a 6 , since no nonvoid member of E can precede a 6 . Hence, the language of canonical representatives compatible with the spines (l; f1g) is In particular, (s; t; t are the members compatible with the spine (2; f1g), so c 5. Hence, using the decomposition of K 6 determined above, we obtain The canonical representatives of the central commutativity classes compatible with spines of the form (l; must have a factorization in which there are at least three occurrences of the words a i , the first and penultimate of these being a 5 or a 6 . Since a 6 cannot be preceded by any of the factors a i , a 5 must be the penultimate factor. Since a 5 can be followed only by a 1 , the first factor must therefore be a 6 , there is no non-void member of E preceding a 6 , and the last factor must be a 1 . From the above decompositions of K 6 and K that the language formed by the members of L that start with a 6 and terminate with a 5 a 1 is a 6 \Delta fa 2 a 4 ; a 2 a 4 a 3 ; a 2 a 4 a 3 a 4 g \Delta a 2 a 4 a 3 a 5 a and therefore Combining our expressions for C i the generating function provided by Theorem 2.6 can be simplified to the form 4. Fully Commutative Involutions We will say that a commutativity class C is palindromic if it includes the reverse of some (equivalently, all) of its members. A fully commutative w 2 W is an involution if and only if R(w) is palindromic. In the following, we will adopt the convention that if X is a set of commutativity classes, then - X denotes the set of palindromic members of X. Similarly, - W and - W FC shall denote the set of involutions in W and W FC , respectively. 4.1 The generic generating function. Consider the enumeration of fully commutative involutions in a series of Coxeter groups of the type considered in Section 2. It is clear that w 2 W FC n is an involution if and only if its branch and central portions are palindromic. Thus by Lemma 2.1, determining the cardinality of W FC n can be split into two subproblems: enumerating - Bn (oe) (the palindromic branch classes) and - CW (oe) (the palindromic central classes). For integers n; l - 0, we define - e. Lemma 4.1. We have Bn (l; ?) Bn (l; f1g) Proof. Following the proof of Lemma 2.3, for n;l denote the cardinality of - Bn (oe) for respectively. Recall that the occurrences of s 1 and s 2 must be interlaced in any representative of B 2 Bn (l; ?), and that when we restrict B to fs ng (and shift indices), we obtain a member of denotes the number of occurrences of s 2 . To be palindromic, it is therefore necessary and sufficient that the fs ng-restriction of B is palindromic, and that l (or 0, if l = 0). This yields the recurrence It is easy to verify that - satisfies the same recurrence and initial conditions. For spines of the form oe = (l; f1g), it is clear that there can be no palindromic classes since for l ? 2, there must be an occurrence of s 2 between the last two occurrences of s 1 , but not for the first two. Assuming l = 2, the bijection provided in the proof of Lemma 2.3 preserves palindromicity, and thus proves the recurrence It is routine to check that the claimed formula for - n;2 satisfies the same recurrence and initial conditions. For the left and right sides of any palindromic B 2 Bn (oe) must be both open or both closed, in the sense defined in the proof of Lemma 2.3. Furthermore, a branch class with this property is palindromic if and only if its restriction to fs ng is palindromic, so the bijection provided in Lemma 2.3 for this case yields and - n;2 . Once again, it is routine to check that the claimed formula for - n;l satisfies the same recurrence and initial conditions. Lemma 4.2. We have Proof. We have We can interpret F l;j (x) as the generating function for sequences of votes in an election in which A defeats B by l votes. Such sequences can be uniquely factored by cutting the sequence after the last moment when B trails A by i votes, 1. The first factor consists of an arbitrary sequence for a tie vote, which has generating function 1= and the remaining l factors each consist of a vote for A, followed by a "ballot sequence" for a tie vote (cf. Section 2.3), which has generating function xR(x 2 ). Turning now to the palindromic central commutativity classes, let us define and associated generating functions l-0 l-3 Theorem 4.3. If W is one of the Coxeter groups displayed in Figure 5, then Proof. As noted previously, w 2 W FC is an involution if and only if the central and branch portions of w are palindromic. Successive applications of Lemmas 2.1, 2.2, and 4.1 therefore yield oe Bn (oe) l-0 l-3 n;l l-0 l-3 The corresponding generating function thus takes the form l-0 l-3 -c l;2 using Lemma 4.2. As we shall see below, both - are rational, so the generating series for belongs to the algebraic function field Q(x; R(x 2 4.2 The A-series. In this case, we have - since there are only two central commutativity classes (namely, those of ( ) and (s)), and both are palindromic. Hence A FC Either by extracting the coefficient of x n , or more directly from (4.1), we obtain A FC 4.3 The B-series. In this case, the central commutativity classes are singletons in which the occurrences of s and t alternate. For each l - 0, there are two such words that are palindromic and have l occurrences of s. Among these, (s; t; s) is the only one that is compatible with a spine (l; A) with A 6= ?. Hence - Theorem 4.3 implies Extracting the coefficient of x dn=2e 4.4 The D-series. In this case, the palindromic central classes are represented by the odd-length subwords of (s; t; s; t middle term is t or t 0 , together with ( ), (s), (t; t 0 ), and In particular, leaving aside (s; t; t there are exactly four such words with l occurrences of s for each even l - 0, so we have Also - is the only representative compatible with a spine of the form (l; with A 6= ?. After simplifying the expression in Theorem 4.3, we obtain Extracting the coefficient of x n\Gamma2 yields (n+1)=2 4.5 The H-series. The palindromic central classes in this case are the same as those for the B-series; the only difference is that those corresponding to hs; ti 7 are now compatible with spines of the form (l; f1; l \Gamma 1g) for l - 4. Thus we have The generating function provided by Theorem 4.3 is therefore and hence fi fi - 4.6 The F -series. Recall that in Section 3.5, we selected a set of canonical representatives for the central commutativity classes by forbidding the subword (s; u). If s is one such representative, let s denote the canonical representative obtained by reversing s and then reversing each offending (s; u)-subword. If s is the canonical representative of a palindromic class (i.e., else s has a unique factorization fitting one of the forms a where a is itself a canonical representative for some central commutativity class. Con- versely, any canonical representative ending in (s) can be uniquely factored into one of the two forms a \Delta (s) or a \Delta (u; s), and the corresponding word obtained by appending a remains central. Similarly, any canonical representative ending with (t) but not (u; t) or factored into the form a \Delta (t), remains central when a is appended. Now from (3.5), the language of canonical representatives ending in (s) consists of the exceptional set f(u; s); (t; u; s)g, together with and the language of representatives ending with (t) but not (u; t) or (u; s; t) is Including the exceptional cases ( ) and (s), this yields The unique palindromic classes compatible with the spines (2; f1g) and (3; f1; 2g) are represented by (s; t; s) and (s; t; u; s; t; s). For the spines recall from Section 3.5 that a canonical representative compatible with oe must begin with (s; t; u; s; t) and end with (t; u; s; t; s). Selecting the portions of (4.6) and (4.7) that begin with (s; t; u; s; t) yields the languages so we have Simplification of the generating series provided by Theorem 4.3 yields F FC The coefficients can be expressed in terms of the Fibonacci numbers as follows: F FC F FC 4.7 The E-series. In Section 3.6, we selected a canonical representative s for each central commutativity class. As in the previous section, we let s denote the canonical representative for the commutativity class of the reverse of s. If s 2 S is a representative of any palindromic commutativity class, then the set of generators appearing an odd number of times in s must commute pairwise. Indeed, the "middle" occurrence of one generator would otherwise precede the "middle" occurrence of some other generator in every member of the commutativity class. Aside from the exceptional cases ( ), (u), and (which cannot be followed and preceded by the same member of S and remain central), it follows that every central palindromic class has a unique representative fitting one of the forms a (t) a; a where a is the canonical representative of some central commutativity class. However, we cannot assert that the above representatives are themselves canonical; for example, if then a (u; t 0 ) a is a representative of a central palindromic class, but the canonical representative of this class is (t; u; s; t For the representatives whose middle factor is s must be the first term of a, assuming that a is nonvoid. Furthermore, if we prepend an initial s (or s; t, in the case of (u; t 0 )), the resulting words (s; t) a, (s; and (s; t; u; t 0 ) a are (in the notation of Section 3.6) members of the formal languages respectively. Conversely, any member of these languages arises in this fashion. For a representative whose middle factor is (u; s), if we prepend (s; t; t 0 ) to (u; s) a, we obtain a member of a central commutativity class whose canonical representative is hence a member of K 6 f(); (s)g. Conversely, any member of K 6 f(); (s)g other than a way. Collecting the contributions of the five types of palindromic central classes, along with the exceptional cases f( ); (u); (s)g, we obtain For the spine there is a unique oe-compatible central class that is palin- dromic; namely, the class of (s; t; t 0 ; s). For the spines oe of the form (l; f1; l \Gamma 1g), recall from (3.10) that the canonical representatives of the oe-compatible classes all begin with a 6 a 2 a 4 and end with a 3 a 5 a 1 . It follows that for a palindromic central class represented by a word of the form (4.9) to be compatible with oe, it is necessary and sufficient that a end with a 3 a 5 a 1 . Using the decompositions obtained in Section 3.6, we find that a 4 a 3 ; a 2 a 4 a 3 a 4 g \Delta a 2 a 4 a 3 a 5 ; a 4 a 3 ; a 2 a 4 a 3 a 4 g \Delta a 2 a 4 a 3 a 5 ; a 6 \Delta fa 2 a 4 ; a 2 a 4 a 3 ; a 2 a 4 a 3 a 4 g \Delta a 2 a 4 a 3 a 5 are the respective portions of K 2 , K 4 , and K 6 that end with a 3 a 5 ; there are no such words in K 5 . It follows that The generating function provided by Theorem 4.3 can be simplified to the form 5. Asymptotics Given the lack of simple expressions for the number of fully commutative members of En and Fn , it is natural to consider asymptotic formulas. Theorem 5.1. We have (a) 1:466 is the real root of x (b) 1:618 is the largest root of x Proof. Consider the generating function Theorem 2.6. In the case of Fn , we see from (3.7) that the singularities of G(x) consist of a branch cut at together with simple poles at and the zeroes of . The latter are (respectively) f1=fl; \Gammafl g, 5)=2 denotes the golden ratio. The smallest of these (in absolute value) is 0:236, a zero of In particular, since 1=fl 3 ! 1=4, the asymptotic behavior of fi fi is governed by the local behavior of G(x) at specifically, since there is a simple pole at using (3.7), together with the relations In the case of En , we see from (3.11) that the singularities of G(x) consist of a branch cut at together with simple poles at and the zeroes of These polynomials are related by the fact that if ff is any zero of is the minimal polynomial of is the minimal polynomial of ff=(1+ ff) 2 . (The fact that such a simple relationship exists is not coincidental; see Remark 5.3 below.) The smallest of the nine zeroes of these polynomials (in absolute value) is 0:682 is the real zero of Equivalently, we have 1=ff is the real root of x 1=4, the asymptotic behavior of is once again governed by the local behavior of G(x) near a simple pole. In this case, we obtain using (3.11) and the fact that Remark 5.2. For the sake of completeness, it is natural also to consider the asymptotic number of fully commutative elements in An , Bn , Dn , and Hn . Given the explicit formulas (3.1), (3.2), (3.3), and (3.4), it is easily established that using Stirling's formula. In each of these cases, the dominant singularity in the corresponding generating function is the branch cut at Remark 5.3. If ff is a pole of f(x), then ff=(1 + ff) is a pole of f(x=(1 \Gamma x)) and ff=(1+ff) 2 is a pole (of some branch) of f(R(x) \Gamma 1). On the other hand, from Theorem 2.6, we see that aside from the branch cut at a pole at the singularities of are limited to those of C 2 (x), C Thus, unless there is unexpected cancellation, for each pole ff of C 2 (x), there will be a triple of poles at ff=(1 in G(x). Now consider the asymptotic enumeration of fully commutative involutions. Again, given the explicit formulas (4.2), (4.3), (4.4), and (4.5), it is routine to show that A FC In the following, fi and fl retain their meanings from Theorem 5.1. Theorem 5.4. We have (a) (b) (c) F FC (d) F FC Proof. Consider the generating series - Theorem 4.3. In the case of Fn , we see from (4.8) that the singularities of - G(x) consist of branch cuts at together with simple poles at (the zeroes of (the zeroes of In absolute value, the smallest of these occur at that the asymptotic behavior of F FC fi fi is determined by the local behavior of - G(x) at specifically, we have F FC F FC Using (4.8) and the fact that Q(fl \Gamma3=2 and a similar calculation (details omitted) yields In the case of En , we see from (4.10) that the singularities of - G(x) consist of branch cuts at together with simple poles at and the square roots of the zeroes of Continuing the notation from the proof of Theorem 5.1, the poles occurring closest to the origin are at 2). Thus we have Using (4.10) and the fact that Q(ffi 1=2 and a similar calculation can be used to determine c \Gamma ; we omit the details. Appendix An Bn Dn En Fn Hn 9 16796 53481 29171 44199 153584 182720 Table 1: The number of fully commutative elements. 2 An Bn Dn En Fn Hn 9 252 637 381 443 968 1014 Table 2: The number of fully commutative involutions. 2 The parenthetical entries correspond to cases in which the group in question is either reducible or isomorphic to a group listed elsewhere. --R Some combinatorial properties of Schubert polynomials "Groupes et Alg'ebres de Lie, Chp. IV-VI," "Advanced Combinatorics," "A Hecke Algebra Quotient and Properties of Commutative Elements of a Weyl Group," Structure of a Hecke algebra quotient "Combinatorial Enumeration," "Modular Representations of Hecke Algebras and Related Algebras," "Reflection Groups and Coxeter Groups," On the fully commutative elements of Coxeter groups Some combinatorial aspects of reduced words in finite Coxeter groups --TR --CTR Sara C. Billey , Gregory S. Warrington, Kazhdan-Lusztig Polynomials for 321-Hexagon-Avoiding Permutations, Journal of Algebraic Combinatorics: An International Journal, v.13 n.2, p.111-136, March 2001

coxeter group;braid relation;reduced word

292363

An Iterative Perturbation Method for the Pressure Equation in the Simulation of Miscible Displacement in Porous Media.

The miscible displacement problem in porous media is modeled by a nonlinear coupled system of two partial differential equations: the pressure-velocity equation and the concentration equation. An iterative perturbation procedure is proposed and analyzed for the pressure-velocity equation, which is capable of producing as accurate a velocity approximation as the mixed finite element method, and which requires the solution of symmetric positive definite linear systems. Only the velocity variable is involved in the linear systems, and the pressure variable is obtained by substitution. Trivially applying perturbation methods can only give an error $O(\epsilon)$, while our iterative scheme can improve the error to $O(\epsilon^m)$ at the $m$th iteration level, where $\epsilon$ is a small positive number. Thus the convergence rate of our iterative procedure is $O(\epsilon)$, and consequently a small number of iterations is required. Theoretical convergence analysis and numerical experiments are presented to show the efficiency and accuracy of our method.

Introduction Miscible displacement occurs, for example, in the tertiary oil-recovery process which can enhance hydrocarbon recovery in the petroleum reservoir. This process involves the injection of a solvent at injection wells with the intention of displacing resident oil to production wells. The resident oil may have been left behind after primary production by reservoir pressure and secondary production by waterflooding. Since the tertiary process requires expensive chemicals and the performance of the displacement is not guaranteed, its numerical simulation plays an important role in determining whether enough additional oil is recovered to make the expense worthwhile and in optimizing the recovery process of hydrocarbon. Department of Mathematics/Institute of Applied Mathematics, University of British Columbia, Vancouver, British Columbia, Canada V6T 1Z2. E-mail address: [email protected] y Department of Mathematics, Wayne State University, Detroit, Michigan 48202, USA. E-mail address: [email protected] or [email protected] Mathematically, miscible displacement in porous media is modeled by a nonlinear coupled system of two partial differential equations with appropriate boundary and initial conditions. The pressure-velocity equation is elliptic, while the concentration equation is parabolic but normally convection-dominated. The concentration equation is derived from the conservation of mass which involves the Darcy velocity of the fluid mixture, but the pressure variable does not appear in the concentration equation. Thus a good approximation of the concentration equation requires accurate solution for the velocity variable. Mixed finite element methods have been applied [9, 10, 11, 15, 16, 25] to the pressure-velocity equation, which can yield the velocity one order more accurate than the finite difference or element methods. However, the finite-dimensional spaces for velocity and pressure need to satisfy the Babuska-Brezzi condition, and the resulting linear system has a nonpositive definite coefficient matrix. Besides, the number of degrees of freedom in the linear system doubles that of finite difference or element methods. Thus efforts have been made to improve the performance of mixed finite element methods [5, 18, 12]. The purpose of this paper is to propose and analyze an iterative perturbation method for the pressure-velocity equation, which is capable of producing as accurate velocity approximation as the mixed finite element method, and requires the solution of symmetric positive definite linear systems. Only the velocity variable is involved in the linear systems and the pressure variable is obtained by substitution without solving any linear systems at each iteration level. Aside from this, the finite element spaces for our velocity and pressure variables need not satisfy the so-called Babuska-Brezzi condition. The iterative perturbation method is a variant of the augmented Lagrangian method [18, 5, 6] applied directly to the continuous differential problem, which can also be viewed as the sequential regularization method [2, 20, 3] applied to stationary problems. Unlike the augmented Lagrangian method using spectral analysis to discuss the convergence and its rate for the discretized problems, we use the method of asymptotic expansion directly for the differential problem following the idea in the sequential regularization method [2, 20]. The asymptotic method is easier to use for more general and more complicated problems than spectral analysis since the former applies readily to non-symmetric and infinite-dimensional operators. Trivially applying perturbation methods (e.g. penalty methods) can only give an error O(ffl), where ffl is a small positive perturbation parameter. We will prove that our iterative scheme can improve the error to O(ffl m ) at the m-th iteration level. In other words, the convergence rate of our iterative procedure is O(ffl). Theoretical convergence analysis and numerical experiments show that the number of iterations is extremely small, usually two or three. The organization of the rest of the paper is as follows. In Section 2, we describe our iterative perturbation method for the time-discretized problem at the partial differential equation level. In Section 3 we show its convergence and the rate of convergence. Then in Section 4 we give a fully-discretized version for the iterative perturbation method for the pressure-velocity equation and a Galerkin method for the concentration equation. Finally, in Section 5 we present numerical examples to demonstrate the effectiveness and accuracy of our method, and in Section 6 we give some remarks and future directions. 2 The Iterative Perturbation Method Consider the miscible displacement of one incompressible fluid by another in a porous reservoir \Omega ae R 2 over a time period denote the pressure and Darcy velocity of the fluid mixture, and c the concentration of the invading fluid. Then the mathematical model is a coupled nonlinear system of partial differential equations 2\Omega \Theta J; (1) div 2\Omega \Theta J; (2) OE @c @t 2\Omega \Theta J; with the boundary conditions and initial condition c) is the mobility tensor of the fluid mixture, fl(x; c) and d(x) are the gravity and vertical coordinate, q is the imposed external rates of flow, OE(x) is the porosity of the rock, D(x; t) is the coefficient of molecular diffusion and mechanical dispersion of one fluid into the other, c) is a known linear function of c representing sources, and - is the exterior normal to the boundary We assume that the mobility is symmetric and positive definite. For existence of p, we assume that the mean value of q is zero and for uniqueness we impose that p have mean value zero. In recent years much attention has been devoted to the numerical simulation of this problem. In this paper we are interested in solving the velocity-pressure equation (1)-(2) using an iterative perturbation method (IPM) for the time-discretized problem. The iterative perturbation method considered herein is a variant of augmented Lagrangian methods [18, 5] and sequential regularization methods [2, 20]. We will analyze the IPM using the technique presented in [2, 20], since the analysis gives the convergence of the iterative procedure and its convergence rate at the same time, and is applicable to non-symmetric problems. After a time discretization, we obtain the following system for u and p at current time step: 2\Omega \Theta J; div 2\Omega \Theta J; where ~ c is an approximation of c assumed to be known here. Taking ffl to be a small positive number, we replace the system (1), (2), (4) by the following iterative perturbation method such that a(~c) 2\Omega \Theta J; and 2\Omega \Theta J; where the initial guess p 0 is required to satisfy the zero mean value property: R p s has mean value zero from (11) and (10). We note that, if taking p 0 j 0, each iteration is a kind of penalty method (probably without its meaning in the optimization context). Our iterative perturbation method (IPM) has the following salient features: 1. We solve a small system (9)-(10) for velocity u, and get pressure p from (11) directly. We will show that the accuracy of such a method is O(ffl s ) at the s-th iteration level. Note that the system (9)-(10) is well-posed since, unlike the usual penalty method, we need not take ffl very small. 2. The velocity-pressure equation was recently solved by the mixed finite element method [9, 10, 11, 15, 16, 25], in which the resulting linear system has a non-positive definite coefficient matrix, and the discrete spaces for u and p need to satisfy the Babuska-Brezzi condition. While in the IPM procedure, the system (9)-(10) leads to a symmetric positive definite coefficient matrix. Since u and p are obtained from equations (9)-(10) and (11) separately, compatibility conditions between the discrete spaces for u and p are not needed. 3. When the standard finite element method [7, 8, 13, 14, 17, 23, 24] is applied for the pressure equation, the velocity need to be obtained by finite differencing the pressure variable, which gives less accuracy. The velocity in our method is obtained directly without finite differencing. Note that the accuracy of the approximate velocity is important, since the concentration equation involves the velocity only. 4. We will see that in xx4 and 5 the discrete version of our IPM scheme (9)-(11) gives the same accuracy for the velocity as the mixed method and requires the solution of well-conditioned linear systems like Galerkin methods. Note that our numerical experiments will show that a few (usually two or three) iterations are enough for the IPM scheme. We now describe some notation to be used throughout the rest of the paper. As usual, we use kuk p and kukH m to denote the standard norms in the Sobolev spaces L respectively, where 1 the subscript p in the norm notation when 2. In addition, we define the divergence space with the following two norms: We shall denote by (\Delta; \Delta) and h\Delta; \Deltai the inner product and duality in\Omega and on \Gamma, respectively. For a normed linear space B with norm k \Delta kB and a sufficiently regular function we define ff 2 and kgk L 1 If [ff; we simplify the notation as kgk respectively. We shall also denote by M and K generic constants, which may be different at different occurrences. We use M for continuous problems and K for discrete problems. 3 Convergence Analysis Before stating our convergence theorem, we first give two lemmas. There exists a unique solution fu; pg, where u 2 H(div) and p 2 H 1(\Omega\Gamma , to the problem 2\Omega \Theta J; div 2\Omega \Theta J; where p and q have mean value zero. Furthermore, there exits a constant M such that the following estimate holds: Proof: Substituting (12) into (13) and (14) we see that p satisfies a Poisson equation with Neumann boundary condition a @p @- Noting the zero mean value of p and using the standard results of the Poisson equation, we obtain the uniqueness and existence of p and the estimate An application of the trace inequality[19] leads to the inequality (15) for p. Then the existence, uniqueness and the estimate (15) for u directly. Writing (9)-(11) as 2\Omega \Theta J; div 2\Omega \Theta J; we can obtain a similar estimate for u s and p s from Lemma 1: Lemma 2 For the solution of the following system: a(~c) 2\Omega \Theta J; we have the stability estimate: Proof: Multiplying both sides of the first equation by u and applying Green's formula leads to (20). We are now ready to describe our convergence theorem and prove it using the method of asymptotic expansion. Theorem 1 Let fu; pg be the solution of system (7), (8), and (4) and fu s the solution of (9)-(11). Then there exists a constant M , independent of s and ffl, such that Proof: We first consider the case in (9). Comparing the coefficients of like powers of ffl, we thus have 5div 5div where the boundary condition (10). The equation (22) has infinitely many solutions in general. We should choose u 10 not only to satisfy (22), but also to ensure that the solution of (23) exists. A choice of u 10 is the exact solution u of (7), (8), and (4), i.e. where p is the exact solution of (7), (8), and (4), and satisfying zero mean value condition. Hence, from Lemma 1, we have has the following form 5 div u The equation (28) suggests that we choose u 11 and a corresponding p 11 (with zero mean value) to satisfy div According to Lemma 1, u 11 and p 11 exist and have the bound: Generally, assuming that u 1(i\Gamma1) and p 1(i\Gamma1) have been found for i - 2, we choose u 1i and p 1i (with zero mean value) satisfying div Applying Lemma 1 and (32), we obtain that all u 1i and p 1i exist uniquely and Next we estimate the remainder of the asymptotic expansion to the m\Gammath power of ffl. Denote the partial sum and the reminder by - Then, from (25)-(27), (29)-(31), and (33)-(35), w 1m satisfies, for m - 1, Then, using (32), (36) and Lemma 2, we obtain Noting that using (36), we thus have proved the estimate (21) for u 1 . On the other hand, (11) can be rewritten as div Using (26), (30), (34) and (39), we have . By taking m - 2, this proves (21) for Now we look at the second iteration Note that (40) gives us a series expansion for p 1 . Plugging the expansions of u 2 and p 1 into and comparing like powers of ffl we obtain div 5 div u 5 div u Again, required to satisfy the boundary condition (10). As in the case of choose u 20 j u and thus have 5 div u This suggests that we construct u 21 and p 21 to satisfy div where p 21 has mean value zero. Obviously u is the solution of (44)-(46). In general, similar to the case of (with zero mean value) to satisfy div m. By the same procedure as in the case of we obtain the error equations similar to (37) and (38) with an addition of a remainder term 5(p 1m ) on the right-hand side. simulating the proof of Lemma 2 and using (40) (i.e. kp we have where the remainder w 2m satisfies the same estimate (39) as w 1m . Noting u 20 j u and u we obtain (21) for u 2 . Then, using (11), (40), (45), (48) and the estimate of w 2m , we conclude that . By taking m - 3, this gives (21) for p 2 since p 21 j 0. We can repeat this procedure, and by induction, conclude the results of the theorem. Remark 2 It is not difficult to see that are a solution to a problem in the form of (12)-(14). Then using the estimates (15) we may obtain a better error estimates for p: From Theorem 1 we see that the convergence rate of our iterative scheme (9)-(11) is O(M ffl) M . This implies that the number of iterations needed to achieve a prescribed accuracy is very small. The fast convergence of our method makes it dramatically different from penalty methods. 4 The Galerkin Approximation In this section, we approximate the velocity-pressure iterative perturbation scheme (9)-(10) and the concentration equation (3) by using the standard Galerkin method. 4.1 The Approximation Scheme \Gammag and Wg. The variational form of (9)-(10) can be written into the following: find such that (a The weak form of the concentration equation (3) reads: find 1(\Omega\Gamma such that @t For h u ? 0 and an integer k - 0, with respect to the velocity-pressure equation, we introduce finite element spaces W h ae W and Y associated with a quasi-regular triangulation of\Omega into triangles or rectangles of diameter less than h u . Similarly, we denote by Z h ae H 1 (\Omega\Gamma the finite-dimensional space for the concentration equation with the grid size h c and approximation index l. Assume that the following approximation properties hold: z h 2Z h where K is a constant. The space W h can be taken to be the vector part of the Raviart-Thomas [22] space of index k, or Brezzi-Douglas-Marini [4] space of index k + 1. Given a partition of J; g. Let fp be fp; u; cg and its approximation at time level t n . We define our approximation scheme at time t n by the following. Step 1: Given C n , find fU follows. Take the initial guess - iteratively - h such that (a where q h is a smoothed source and sink function. Let U U s and P Step 2: When U n known, find C n+1 2 Z h such that \Deltat n Note that in step 1 of the scheme, the initial guess could be more efficiently taken as - Our numerical experiments will show that the number of iterations can usually be taken to be or 3 for a range of perturbation parameter to 10 \Gamma5 . For convection-dominated displacement problems, sharp fluid interfaces move along characteristic or near-characteristic directions. Thus the modified method of characteristics [8, 11, 15, 16, 23, 24, 26] may be applied to treat the concentration equation. In this case, the accurate Darcy velocity computed using the IPM helps to determine the characteristic direction more precisely. 4.2 For our error analysis below we shall make use of the following elliptic projections of fu; cg : (a (div (R u such that (D(u)r(R and q is the right-hand side function of equation (2) [9]. This - function is chosen to assure coercivity of the projection form. Then, it can be shown that [9, 10] ck ck From (56) and the definition of Y and Y h , for Y we can also find fR We now start our error analysis by combining (52), (58) and (61) to get (a Then we see that In view of (63), (65) and (68) we have Presuming that kp u , we thus have (for On the other hand, following the way of getting (67) and combining (52), (11), (58) and (59), we get (a It is not difficult to see that p s presuming p Hence we can find R p satisfying (66). Since - h we can take w 0 2 W h such that div w according to [1, Lemma 3.3] (assuming\Omega is a bounded, Lipschitz-continuous, and connected domain), we have Hence we take in (71) and obtain In view of (66), (69) and (72), we have Thus the error estimates (70) and (74) are reduced to bounding C \Gamma Rc. We denote that Combining (53), (60), (62) and taking z = e n+1 as a test function we obtain the error equation \Deltat n re n+1 ) \Deltat n \Deltat n re We now estimate the right-hand side terms of the error equation (75). For T 1 we have by Since c \Deltat n kc t k 2 we see that \Deltat Assume that D(u) is Lipschitz [9], by (63), (64), (65), and (70) we have c kck 2 where j is a small positive number. Obviously, We rewrite T 5 in the following form that re is bounded, from (63), (64), (65) and (70) we obtain c kck 2 Note that g is a linear function of c, we have c kck 2 Finally, we see that the left-hand side of (75) dominates2\Deltat n positive number. Substituting (76)-(82) into (75) and choosing j small we obtain the following error inequality \Deltat where c kck 2 Applying Gronwall's lemma to (84) we see that \Deltat n We are now ready to state our main results for scheme (58)-(60). Theorem 2 Let fp; u; cg be the solution to Problem (3)-(2), and fP; U; Cg the solution to the scheme (58)-(60) with s iterations at each time step. Then there exists a constant K, for \Deltat sufficiently small, c kck L \Deltat n kC c kck L This theorem tells us that for sufficiently small perturbation parameter ffl, the error estimates for the velocity, pressure and concentration are optimal. U s R U s U s Figure 1. One element with velocity on each edge and pressure at the center. 5 Numerical Experiments In this section, we present some numerical examples to show how well our iterative scheme performs, and how the parameter ffl affects the number of iterations and accuracy. For simplicity, we will just consider the pressure-velocity equation, since the concentration equation has been analyzed previously [9, 10, 11, 14, 15, 16, 17, 23, 24, 25]. Consider the elliptic problem with Neumann boundary condition div where \Omega is a square and \Gamma its boundary. More general domain\Omega will not present technical problems. For simplicity, we also assume that the coefficient a is a scalar. The approximation scheme takes the form: Find U s 2 W h , for (a Partition the domain\Omega into a set of squares of side length h. We take the space W h to be the vector part of the Raviart-Thomas [22] space of index 0. Thus where P k is the set of one variable polynomials of order less than or equal to k. Consequently, the approximate pressure P s lies in the space of piecewise constants. Partitioning the domain into triangles or rectangles, or applying higher order approximation polynomials can be treated analogously. Let U s ff denote the constant value of the flux in positive x or y-direction on the edge ff; L, R, B, T (representing left, right, bottom, and top, respectively), of each element. See Figure 1. Consider w to be the basis function (on the standard reference square). Applying the trapezoidal rule to (88) we where q ff is the value of q at the middle point of edge ff. Similarly, letting simplifying we have the following linear system for each element. R U s Note the equation (89) has the discrete version on each element: From equations (91)-(94) we can easily form the element stiffness matrix. Then assembling all the element matrices and taking into account the boundary condition we obtain the stiffness matrix. The force vector can be gotten in an analogous way. To test the convergence, the following stopping criterion will be adopted where k is a positive integer, and k \Delta k1 denotes the discrete L 1 norm. When this criterion is satisfied, the iterative process is stopped and the solutions at iteration s are adopted. We now apply the iterative procedure (91)-(95) to some test problems on Sun SPARCstation IPC with computation in C++ data type double. We will use the stopping criterion (96) with pressure errors are measured for iterates fU against the exact solutions under the L 1 norm, although in the stopping criterion (96) errors are measured for iterates only. The initial guesses are always chosen to be zero, so all errors are 1.00 before the iterative procedure starts. Example 1: Let velocity u and pressure p satisfy div @ The true solutions for velocity u and pressure p are given by sin(-x) sin(-y) The pressure p and external flow rate div u are chosen in such a way that they both have mean value zero. Example 2: Let velocity u and pressure p satisfy the problem div 1. The true solutions for velocity u and pressure p are given by Example 3: Let velocity u and pressure p satisfy the nonhomogeneous problem div y 10. The function g is chosen such that the true solutions for velocity u and pressure p are given by For Example 1, the results with (uniform) grid sizes 1and 1are shown in Tables 1 through 4. Although more than 2 iterations are required for the iterative procedure to stop when or 10 \Gamma2 , the approximate velocity and pressure at iteration 2 are accurate enough. For Example 2, the results with grid sizes 1and 1are shown in Tables 5 and 6. The results of Example 3 with grid sizes 1and 1are shown in Tables 7 and 8. However, the pressure is less accurate than velocity in this example. This might be caused by the fact errors are calculated relative to the exact solutions and shown in the L 1 norm. Initial guesses are always chosen to be zero, so all relative errors are 1.00 at iteration 0. iteration velocity pressure velocity pressure velocity pressure Table 2: Numerical results for Example 1 with grid size = 1. Both velocity errors and pressure errors are calculated relative to the exact solutions and shown in the L 1 norm. All runs stop after 2 iterations except for the case \Gamma6 in which 1 iteration is required. iteration velocity pressure velocity pressure velocity pressure that the pressure iterates fP s g converge to the true pressure p up to a constant, since p has mean value zero while fP s g do not. From Tables through 8 we conclude that our iterative method performs as well as the theory predicts. In particular, it can achieve good accuracy for velocity, while the linear systems solved are symmetric and positive definite. Also, the computational work of our method is much smaller than that of mixed methods, since the number of iterations required is usually very small. Table 3: Numerical results for Example 1 with grid size = 1. Both velocity errors and pressure errors are calculated relative to the exact solutions and shown in the L 1 norm. iteration velocity pressure velocity pressure velocity pressure errors are calculated relative to the exact solutions and shown in the L 1 norm. All runs stop after 2 iterations except for the case \Gamma6 in which 1 iteration is required. iteration velocity pressure velocity pressure velocity pressure Table 5: Numerical results for Example 2 with grid size = 1. Both velocity errors and pressure errors are calculated relative to the exact solutions and shown in the L 1 norm. iteration velocity pressure velocity pressure velocity pressure Table Numerical results for Example 2 with grid size = 1. Both velocity errors and pressure errors are calculated relative to the exact solutions and shown in the L 1 norm. iteration velocity pressure velocity pressure velocity pressure Table 7: Numerical results for Example 3 with grid size = 1. Both velocity errors and pressure errors are calculated relative to the exact solutions and shown in the L 1 norm. iteration velocity pressure velocity pressure velocity pressure iteration velocity pressure velocity pressure velocity pressure 6 Concluding Remarks We have proposed an iterative procedure for the pressure-velocity equation in the numerical simulation of miscible displacement in porous media. This procedure is first analyzed at the differential level and then discretized by finite element methods. Theoretical analysis and numerical experiments show that this procedure converges at the rate of O(ffl), where ffl is a small positive number. The fast convergence rate and the ease of choosing relaxation parameter make our iterative method different from penalty methods and Uzawa's algorithm. Indeed, our numerical experiments show that two or three iterations are usually enough for a variety of problems, corresponding to Compared with mixed finite element methods, the discrete version of our scheme can provide the same accurate approximations for velocity and pressure, which is crucial in reservoir problems since velocity is intimately involved in the concentration equation. However, in contrast to mixed finite element methods, our scheme requires only the solution of symmetric and positive definite linear systems which have a smaller number of degrees of freedom corresponding to the velocity variable. Since our method can completely decouple the velocity and pressure variables, the so-called Babuska-Brezzi condition is not needed in constructing the finite dimensional spaces for velocity and pressure. In view of the advantages of our iterative method, we can conclude that it can lead to great savings in computer memory and small execution time of the numerical algorithm. Also, it has the capability of effectively dealing with heterogeneous and anisotropic media in which the permeability tensor may be non-diagonal, rapidly-varying and even discontinuous. In [27], one of the authors conducted numerical experiments for the miscible misplacement using the IPM for the pressure-velocity equation and a modified characteristic method for the concentration equation. Finally, we point out that the iterative procedure presented in this paper can be applied easily to three-dimensional problems. Acknowledgement : The authors would like to thank Professors Uri Ascher and Jim Dou- glas, Jr. for their inspiration and support in the pursuit of this research. --R Decomposition of vector spaces and application to the Stokes Sequential regularization methods for higher index DAEs with constraint singularities: Linear index-2 case Sequential regularization methods for nonlinear higher index DAEs Mixed and Hybrid Finite Element Methods Discontinuous upwinding and mixed finite elements for two-phase flows in reservoir simulation On the approximation of miscible displacement in porous media by a method of characteristics combined with a mixed method Inexact and preconditioned Uzawa algorithms for saddle point problems The Mathematics of Reservoir Simulation Efficient time-stepping methods for miscible displacement problems in porous media Simulation of miscible displacement using mixed methods and a modified method of characteristics Convergence analysis of an approximation of miscible displacement in porous media by mixed finite elements and a modified method of characteristics Galerkin Methods for miscible displacement problems in porous media Applications to the Numerical Solution of Boundary-Value Problems Finite Element Methods for Navier-Stokes Equations A sequential regularization method for time-dependent incompressible Navier-Stokes equations Regularization methods for differential equations and their numerical solution A mixed finite element method for second order elliptic problems Finite elements with characteristic finite element method for a miscible displacement problem Time stepping along characteristics with incomplete iteration for a Galerkin approximation of miscible displacement in porous media Mixed methods with dynamic finite element spaces for miscible displacement in porous media A characteristic mixed method with dynamic finite element space for convection-dominated diffusion problems Numerical simulation of miscible displacement in porous media using an iterative perturbation algorithm combined with a modified method of characteristics. --TR

miscible displacement;iterative method;perturbation method;galerkin method;flow in porous media

292374

A Sparse Approximate Inverse Preconditioner for Nonsymmetric Linear Systems.

This paper is concerned with a new approach to preconditioning for large, sparse linear systems. A procedure for computing an incomplete factorization of the inverse of a nonsymmetric matrix is developed, and the resulting factorized sparse approximate inverse is used as an explicit preconditioner for conjugate gradient--type methods. Some theoretical properties of the preconditioner are discussed, and numerical experiments on test matrices from the Harwell--Boeing collection and from Tim Davis's collection are presented. Our results indicate that the new preconditioner is cheaper to construct than other approximate inverse preconditioners. Furthermore, the new technique insures convergence rates of the preconditioned iteration which are comparable with those obtained with standard implicit preconditioners.

Introduction . In this paper we consider the solution of nonsingular linear systems of the form (1) where the coefficient matrix A 2 IR n\Thetan is large and sparse. In particular, we are concerned with the development of preconditioners for conjugate gradient-type methods. It is well-known that the rate of convergence of such methods for solving (1) is strongly influenced by the spectral properties of A. It is therefore natural to try to transform the original system into one having the same solution but more favorable spectral properties. A preconditioner is a matrix that can be used to accomplish such a transformation. If G is a nonsingular Dipartimento di Matematica, Universit'a di Bologna, Italy and CERFACS, 42 Ave. G. Coriolis, 31057 Toulouse Cedex, France ([email protected]). This work was supported in part by a grant under the scientific cooperation agreement between the CNR and the Czech Academy of Sciences. y Institute of Computer Science, Academy of Sciences of the Czech Republic, Pod vod'arenskou v-e-z'i 2, ([email protected]). The work of this author was supported in part by grants GA CR No. 201/93/0067 and GA AS CR No. 230401 and by NSF under grant number INT-9218024. Michele Benzi and Miroslav T-uma matrix which approximates A linear system (2) will have the same solution as system (1) but the convergence rate of iterative methods applied to (2) may be much higher. Problem (2) is preconditioned from the left, but right preconditioning is also possible. Preconditioning on the right leads to the transformed linear system Once the solution y of (3) has been obtained, the solution of (1) is given by The choice between left or right preconditioning is often dictated by the choice of the iterative method. It is also possible to use both forms of preconditioning at once (split preconditioning), see [3] for further details. Note that in practice it is not required to compute the matrix product GA (or AG) explicitly, because conjugate gradient-type methods only necessitate the coefficient matrix in the form of matrix-vector multiplies. Therefore, applying the preconditioner within a step of a gradient-type method reduces to computing the action of G on a vector. Loosely speaking, the closer G is to the exact inverse of A, the higher the rate of convergence of iterative methods will be. Choosing yields convergence in one step, but of course constructing such a preconditioner is equivalent to solving the original problem. In practice, the preconditioner G should be easily computed and applied, so that the total time for the preconditioned iteration is less than the time for the unpreconditioned one. Typically, the cost of applying the preconditioner at each iteration of a conjugate gradient-type method should be of the same order as the cost of a matrix-vector multiply involving A. For a sparse A, this implies that the preconditioner should also be sparse with a density of nonzeros roughly of the same order as that of A. Clearly, the effectiveness of a preconditioning strategy is strongly problem and architecture dependent. For instance, a preconditioner which is expensive to compute may become viable if it is to be reused many times, since in this case the initial cost of forming the preconditioner can be amortized over several linear systems. This situation occurs, for in- stance, when dealing with time-dependent or nonlinear problems, whose numerical solution gives rise to long sequences of linear systems having the same coefficient matrix (or a slowly varying one) and different right-hand sides. Furthermore, preconditioners that are very efficient in a scalar computing environment may show poor performance on vector and parallel machines, and conversely. Approximate Inverse Preconditioning 3 A number of preconditioning techniques have been proposed in the literature (see, e.g., [2],[3] and the references therein). While it is generally agreed that the construction of efficient general-purpose preconditioners is not possible, there is still considerable interest in developing methods which will perform well on a wide range of problems and are well-suited for state-of-the-art computer architectures. Here we introduce a new algebraic preconditioner based on an incomplete triangular factorization of A \Gamma1 . This paper is the natural continuation of [8], where the focus was restricted to symmetric positive definite systems and to the preconditioned conjugate gradient method (see also [5],[7]). The paper is organized as follows. In x2 we give a quick overview of implicit and explicit preconditioning techniques, considering the relative advantages as well as the limitations of the two approaches. In x3 we summarize some recent work on the most popular approach to approximate inverse preconditioning, based on Frobenius norm minimization. In x4 we introduce the new incomplete inverse triangular decomposition technique and describe some of its theoretical properties. A graph-theoretical characterization of fill-in in the inverse triangular factorization is presented in x5. In x6 we consider the use of preconditioning on matrices which have been reduced to block triangular form. Implementation details and the results of numerical experiments are discussed in xx7 and 8, and some concluding remarks and indications for future work are given in x9. Our experiments suggest that the new preconditioner is cheaper to construct than preconditioners based on the optimization approach. Moreover, good rates of convergence can be achieved by our preconditioner, comparable with those insured by standard ILU-type techniques. 2. Explicit vs. implicit preconditioning. Most existing preconditioners can be broadly classified as being either of the implicit or of the explicit kind. A preconditioner is implicit if its application, within each step of the chosen iterative method, requires the solution of a linear system. A nonsingular matrix M - A implicitly defines an approximate applying G requires solving a linear system with coefficient matrix M . Of course, M should be chosen so that solving a system with matrix M is easier than solving the original problem (1). Perhaps the most important example is provided by preconditioners based on an Incomplete LU (ILU) decomposition. Here U where L and - U are sparse triangular matrices which approximate the exact L and U factors of A. Applying the preconditioner requires the solution of two sparse triangular systems (the forward and backward solves). Other notable examples of implicit preconditioners include the ILQ, SSOR and ADI preconditioners, see [3]. 4 Michele Benzi and Miroslav T-uma In contrast, with explicit preconditioning a matrix G - A \Gamma1 is known (possibly as the product of sparse matrices) and the preconditioning operation reduces to forming one (or more) matrix-vector product. For instance, many polynomial preconditioners belong to this class [37]. Other explicit preconditioners will be described in the subsequent sections. Implicit preconditioners have been intensively studied, and they have been successfully employed in a number of applications. In spite of this, in the last few years an increasing amount of attention has been devoted to alternative forms of preconditioning, especially of the explicit kind. There have been so far two main reasons for this recent trend. In the first place, shortly after the usage of modern high-performance architectures became widespread, it was realized that straightforward implementation of implicit preconditioning in conjugate gradient-like methods could lead to severe degradation of the performance on the new machines. In particular, the sparse triangular solves involved in ILU-type preconditioning were found to be a serial bottleneck (due to the recursive nature of the computation), thus limiting the effectiveness of this approach on vector and parallel computers. It should be mentioned that considerable effort has been devoted to overcoming this difficulty. As a result, for some architectures and types of problems it is possible to introduce nontrivial parallelism and to achieve reasonably good performance in the triangular solves by means of suitable reordering strategies (see, e.g., [1],[38],[54]). However, the triangular solves remain the most problematic aspect of the computation, both on shared memory [33] and distributed memory [10] computers, and for many problems the efficient application of an implicit preconditioner in a parallel environment still represents a serious challenge. Another drawback of implicit preconditioners of the ILU-type is the possibility of break-downs during the incomplete factorization process, due to the occurrence of zero or exceedingly small pivots. This situation typically arises when dealing with matrices which are strongly unsymmetric and/or indefinite, even if pivoting is applied (see [11],[49]), and in general it may even occur for definite problems unless A exhibits some degree of diagonal dominance. Of course, it is always possible to safeguard the incomplete factorization process so that it always runs to completion, producing a nonsingular preconditioner, but there is also no guarantee that the resulting preconditioner will be of acceptable quality. Fur- thermore, as shown in [23], there are problems for which standard ILU techniques produce unstable incomplete factors, resulting in useless preconditioners. Explicit preconditioning techniques, based on directly approximating A \Gamma1 , have been developed in an attempt to avoid or mitigate such difficulties. Applying an explicit preconditioner only requires sparse matrix-vector products, which should be easier to parallelize Approximate Inverse Preconditioning 5 than the sparse triangular solves, and in some cases the construction of the preconditioner itself is well-suited for parallel implementation. In addition, the construction of an approximate inverse is sometimes possible even if the matrix does not have a stable incomplete LU decomposition. Moreover, we mention that sparse incomplete inverses are often used when constructing approximate Schur complements (pivot blocks) for use in incomplete block factorization and other two-level preconditioners, see [2],[3],[12],[15]. Of course, explicit preconditioners are far from being completely trouble-free. Even if a sparse approximate inverse G is computed, care must be exercised to ensure that G is nonsingular. For nonsymmetric problems, the same matrix G could be a good approximate inverse if used for left preconditioning and a poor one if used for right preconditioning, see [36, p. 96],[45, p. 66],[48]. Furthermore, explicit preconditioners are sometimes not as effective as implicit ones at reducing the number of iterations, in the sense that there are problems for which they require a higher number of nonzeros in order to achieve the same rate of convergence insured by implicit preconditioners. One of the reasons for this limitation is that an explicit preconditioner attempts to approximate A \Gamma1 , which is usually dense, with a sparse matrix. Thus, an explicit preconditioner is more likely to work well if A \Gamma1 contains many entries which are small (in magnitude). A favorable situation is when A exhibits some form of diagonal dominance, but for such problems implicit preconditioning is also likely to be very effective. Hence, for problems of this type, explicit preconditioners can be competitive with implicit ones only if explicitness is fully exploited. Finally, explicit preconditioners are usually more expensive to compute than implicit ones, although this difference may become negligible in the common situation where several linear systems with the same coefficient matrix and different right-hand sides have to be solved. In this case the time for computing the preconditioner is often only a fraction of the time required for the overall computation. It is also worth repeating that the construction of certain sparse approximate inverses can be done, at least in principle, in a highly parallel manner, whereas the scope for parallelism in the construction of ILU-type preconditioners is more limited. 3. Methods based on Frobenius norm minimization. A good deal of work has been devoted to explicit preconditioning based on the following approach: the sparse approximate inverse is computed as the matrix G which minimizes kI \Gamma GAk (or kI \Gamma AGk for right preconditioning) subject to some sparsity constraint (see [4], Ch. 8 of [2],[16],[43], [44],[32],[31],[11],[30]). Here the matrix norm is usually the Frobenius norm or a weighted variant of it, for computational reasons. With this choice, the constrained minimization problem decouples into n independent linear least squares problems (one for each row, or 6 Michele Benzi and Miroslav T-uma column of G), the number of unknowns for each problem being equal to the number of nonzeros allowed in each row (or column) of G. This immediately follows from the identity is the ith unit vector and g i is the ith column of G. Clearly, there is considerable scope for parallelism in this approach. The resulting sparse least squares problems can be solved, in principle, independently of each other, either by direct methods (as in [44], [31],[30]) or iteratively ([11],[42]). In early papers (e.g. [4],[32],[43]) the sparsity constraint was imposed a priori, and the minimizer was found relative to a class of matrices with a predetermined sparsity pattern. For instance, when A is a band matrix with a good degree of diagonal dominance, a banded approximation to A \Gamma1 is justified, see [18]. However, for general sparse matrices it is very difficult to guess a good sparsity pattern for an approximate inverse, and several recent papers have addressed the problem of adaptively defining the nonzero pattern of G in order to capture "large" entries of the inverse [31],[30]. Indeed, by monitoring the size of each residual it is possible to decide whether new entries of g i are to be retained or discarded, in a dynamic fashion. Moreover, the information on the residuals can be utilized to derive rigorous bounds on the clustering of the singular values of the preconditioned matrix and therefore to estimate its condition number [31]. It is also possible to formulate conditions on the norm of the residuals which insure that the approximate inverse will be nonsingular. Unfortunately, such conditions appear to be of dubious practical value, because trying to fulfill them could lead to a very dense approximate inverse [16],[11]. A disadvantage of this approach is that symmetry in the coefficient matrix cannot be exploited. If A is symmetric positive definite (SPD), the sparse approximate inverse will not be symmetric in general. Even if a preset, symmetric sparsity pattern is enforced, there is no guarantee that the approximate inverse will be positive definite. This could lead to a breakdown in the conjugate gradient acceleration. For this reason, Kolotilina and Yeremin [43],[44] propose to compute an explicit preconditioner of the form is lower triangular. The preconditioned matrix is then GLAG T , which is SPD, and the conjugate gradient method can be applied. The matrix GL is the solution of a constrained minimization problem for the Frobenius norm of I \Gamma LGL where L is the Cholesky factor of A. In [43] it is shown how this problem can be solved without explicit knowledge of any of the entries of L, using only entries of the coefficient matrix A. The same technique can also be used to compute a factorized approximate inverse of a nonsymmetric matrix by Approximate Inverse Preconditioning 7 separately approximating the inverses of the L and U factors. As it stands, however, this technique requires that the sparsity pattern of the approximate inverse triangular factors be specified in advance, and therefore is not suitable for matrices with a general sparsity pattern. There are additional reasons for considering factorized approximate inverses. Clearly, with the approximate inverse G expressed as the product of two triangular factors it is trivial to insure that G is nonsingular. Another argument in favor of this approach is given in [11], where it is observed that factorized forms of general sparse matrices contain more information for the same storage than if a single product was stored. The serial cost for the construction of this type of preconditioner is usually very high, although the theoretical parallel complexity can be quite moderate [44],[30]. The results of numerical experiments reported in [44] demonstrate that factorized sparse approximate inverse preconditioners can insure rapid convergence of the preconditioned conjugate gradient iteration when applied to certain finite element discretizations of 3D PDE problems arising in elasticity theory. However, in these experiments the preconditioning strategy is not applied to the coefficient matrix directly, but rather to a reduced system (Schur comple- ment) which is better conditioned and considerably less sparse than the original problem. When the approximate inverse preconditioner is applied directly to the original stiffness matrix A, the rate of convergence of the PCG iteration is rather disappointing. A comparison between a Frobenius norm-based sparse approximate inverse preconditioner and the ILU(0) preconditioner on a number of general sparse matrices has been made in [30]. The reported results show that the explicit preconditioner can insure rates of convergence comparable with those achieved with the implicit ILU-type approach. Again, it is observed that the construction of the approximate inverse is often very costly, but amenable to parallelization. Factorized sparse approximate inverses have also been considered by other authors, for instance by Kaporin [39],[40],[41], whose approach is also based on minimizing a certain matrix functional and is closely related to that of Kolotilina and Yeremin. In the next sections we present an alternative approach to factorized sparse approximate inverse preconditioning which is not grounded in optimization, but is based instead on a direct method of matrix inversion. As we shall see, the serial cost of forming a sparse approximate inverse with this technique is usually much less than with the optimization approach, while the convergence rates are still comparable, on average, with those obtained with ILU-type preconditioning. 8 Michele Benzi and Miroslav T-uma 4. A method based on inverse triangular factorization. The optimization approach to constructing approximate inverses is not the only possible one. In this section we consider an alternative procedure based on a direct method of matrix inversion, performed incompletely in order to preserve sparsity. This results in a factorized sparse G - A \Gamma1 . Being an incomplete matrix factorization method, our procedure resembles classical ILU- type implicit techniques, and indeed we can draw from the experience accumulated in years of use of ILU-type preconditioning both at the implementation stage and when deriving theoretical properties of the preconditioner G. This paper continues the work in [8], where the symmetric positive definite case was studied (see also [5],[7]). The construction of our preconditioner is based on an algorithm which computes two sets of vectors fz i g n , which are A-biconjugate, i.e. such that w T only if i 6= j. Given a nonsingular matrix A 2 IR n\Thetan , there is a close relationship between the problem of inverting A and that of computing two sets of A-biconjugate vectors fz i g n and fw i g n . If is the matrix whose ith column is z i and is the matrix whose ith column is w i , then It follows that W and Z are necessarily nonsingular and Hence, the inverse of A is known if two complete sets of A-biconjugate vectors are known. Note that there are infinitely many such sets. Matrices W and Z whose columns are A- biconjugate can be explicitly computed by means of a biconjugation process applied to the columns of any two nonsingular matrices W (0) , Z (0) 2 IR n\Thetan . A computationally convenient choice is to let W the biconjugation process is applied to the unit basis vectors. In order to describe the procedure, let a T and c T denote the ith row of A and A T , Approximate Inverse Preconditioning 9 respectively (i.e., c i is the ith column of A). The basic A-biconjugation procedure can be written as follows. THE BICONJUGATION ALGORITHM := a T z (i) z (i\Gamma1) and Return This algorithm is essentially due to L. Fox, see Ch. 6 of [25]. Closely related methods have also been considered by Hestenes and Stiefel [35, pp. 426-427],[34] and by Stewart [52]. A more general treatment is given in the recent paper [14]. Geometrically, the procedure can be regarded as a generalized Gram-Schmidt orthogonalization with oblique projections and nonstandard inner products, see [6],[14]. Several observations regarding this algorithm are in order. In the first place we note that the above formulation contains some redundancy, since in exact arithmetic Therefore, at step i the computation of the dot product q (i\Gamma1) i may be replaced by the assignment q (i\Gamma1) . Another observation is the fact that the procedure, as it stands, is vulnerable to breakdown (division by zero), which occurs whenever any of the Michele Benzi and Miroslav T-uma quantities p (i\Gamma1) happens to be zero. It can be shown that in exact arithmetic, the biconjugation algorithm will not break down if and only if all the leading principal minors of A are nonzero (see below). For any nonsingular matrix A there exists a permutation matrix (or Q) such that the procedure applied to PA (or to AQ) will not break down. As in the LU decomposition with pivoting, such permutation matrices represent row (or column) interchanges on A which can be performed, if needed, in the course of the computation. If the biconjugation process can be carried to completion without interchanges, the resulting Z and W matrices are upper triangular, 1 they both have all diagonal entries equal to one, and satisfy the identity We recognize in (5) the familiar LDU decomposition A = LDU , where L is unit lower triangular, U is unit upper triangular and D is the diagonal matrix with the pivots down the main diagonal. Because this factorization is unique, we have that the biconjugation algorithm explicitly computes and the matrix D, which is exactly the same in (5) and in A = LDU . Hence, the process produces an inverse triangular decomposition of A or, equivalently, a triangular decomposition (of the UDL type) of A \Gamma1 . The p i 's returned by the algorithm are the pivots in the LDU factorization of A. If we denote by \Delta i the ith leading principal minor of A (1 - i - n) and let the identity (5) implies that showing that the biconjugation algorithm can be performed without breakdowns if and only if all leading principal minors of A are non-vanishing. In finite precision arithmetic, pivoting may be required to promote numerical stability. Once Z, W and D are available, the solution of a linear system of the form (1) can be computed, by (4), as 1 Note that this is not necessarily true when a matrix other than the identity is used at the outset, i.e. if Approximate Inverse Preconditioning 11 In practice, this direct method for solving linear systems is not used on account of its cost: for a dense n \Theta n matrix, the biconjugation process requires about twice the work as the LU factorization of A. Notice that the cost of the solve phase using (6) is the same as for the forward and backward solves with the L and U factors. If A is symmetric, the number of operations in the biconjugation algorithm can be halved by observing that W must equal Z. Hence, the process can be carried out using only the rows of A, the z-vectors and the p (i\Gamma1) . The columns of the resulting Z form a set of conjugate directions for A. If A is SPD, no breakdown can occur (in exact arithmetic), so that pivoting is not required and the algorithm computes the L T DL factorization of A \Gamma1 . This method was first described in [26]. Geometrically, it amounts to Gram-Schmidt orthogonalization with inner product hx; yi := x T Ay applied to the unit vectors e It is sometimes referred to as the conjugate Gram-Schmidt process . The method is still impractical as a direct solver because it requires about twice the work of Cholesky for dense matrices. However, as explained in [5] and [6], the same algorithm can also be applied to nonsymmetric systems, resulting in an implicit LDU factorization where only are computed. Indeed, it is possible to compute a solution to (1) for any right-hand side using just Z, D and part of the entries of A. This method has the same arithmetic complexity as Gaussian elimination when applied to dense problems. When combined with suitable sparsity-preserving strategies the method can be useful as a sparse direct solver, at least for some types of problems (see [5],[6]). For a sparse symmetric and positive definite A, the Z matrix produced by the algorithm tends to be dense (see the next section), but it can be observed experimentally that very often, most of the entries in Z have small magnitude. If fill-in in the Z matrix is reduced by removing suitably small entries in the computation of the z-vectors, the algorithm computes a sparse matrix - Z and a diagonal matrix - D such that (i.e., a factorized sparse approximate inverse of A). Hence, G can be used as an explicit preconditioner for the conjugate gradient method. A detailed study of this preconditioning strategy for SPD problems can be found in [8], where it is proven that the incomplete inverse factorization exists if A is an H-matrix (analogously to ILU-type factorizations). The numerical experiments in [8] show that this approach can insure fast convergence of the PCG iteration, almost as good as with implicit preconditioning of the incomplete Cholesky type. The construction of the preconditioner itself, while somewhat more expensive than the computation of the incomplete Cholesky factorization, is still quite cheap. This is in contrast Michele Benzi and Miroslav T-uma with the least squares approach described in the previous section, where the construction of the approximate inverse is usually very time consuming, at least in a sequential environment. In the remainder of this paper we consider an explicit preconditioning strategy based on the biconjugation process described above. Sparsity in the Z and W factors of A \Gamma1 is preserved by removing "small" fill in the z- and w-vectors. A possibility would be to drop all newly added fill-in elements outside of a preset sparsity pattern above the main diagonal in Z and W ; however, for general sparse matrices it is very difficult to guess a reasonable sparsity pattern, and a drop tolerance is used instead. A trivial extension of the results in [8] shows that the incomplete biconjugation process (incomplete inverse factorization) cannot break down, in exact arithmetic, if A is an H-matrix. For more general matrices it is necessary to safeguard the computation in order to avoid breakdowns. This requires pivot modifications and perhaps some form of pivoting -we postpone the details to x7. The incomplete biconjugation algorithm computes sparse unit upper triangular matrices W - W and a nonsingular diagonal matrix - D - D such that is a factorized sparse approximate inverse of A which can be used as an explicit preconditioner for conjugate gradient-type methods for the solution of (1). We conclude this section with a few remarks on properties of the approximate inverse preconditioner G just described. If A is not an H-matrix, as already mentioned, the construction of the preconditioner could break down due to the occurrence of zero or extremely small pivots. However, following [46], we note that there always exists ff ? 0 such that ffI is diagonally dominant, and hence an H-matrix. Therefore, if the incomplete bicon- jugation algorithm breaks down, one could try to select ff ? 0 and re-attempt the process on the shifted matrix A should be large enough to insure the existence of the incomplete inverse factorization, but also small enough so that A 0 is close to A. This approach has several drawbacks: for ill-conditioned matrices, the quality of the resulting preconditioner is typically poor; furthermore, the breakdown that prompts the shift may occur near the end of the biconjugation process, and the preconditioner may have to be recomputed several times before a satisfactory value of ff is found. A better strategy is to perform diagonal modifications only as the need arises, shifting pivots away from zero if their magnitude is less than a specified threshold (see x7 for details). If A is an M-matrix, it follows from the results in [8] that G is a nonnegative matrix. Approximate Inverse Preconditioning 13 Moreover, it is easy to see that componentwise the following inequalities hold: where DA is the diagonal part of A. Furthermore, if G 1 and G 2 are two approximate inverses of the M-matrix A produced by the incomplete biconjugation process and the drop tolerance used for G 1 is greater than or equal to the drop tolerance used for G 2 , then The same is true if sparsity patterns are used to determine the nonzero structure in - Z and W and the patterns for G 2 include the patterns for G 1 . This monotonicity property is shared by other sparse approximate inverses, see for example Ch. 8 in [2]. We note that property (7) is important if the approximate inverse is to be used within an incomplete block factorization of an M-matrix A, because it insures that all the intermediate matrices produced in the course of the incomplete factorization preserve the M-matrix property (see [2, pp. 263-264]). Finally, after discussing the similarities, we point to a difference between our incomplete inverse factorization and the ILU-type factorization of a matrix. The incomplete factorization of an M-matrix A induces a splitting which is a regular splitting, and therefore convergent: ae(I \Gamma - denotes the spectral radius of a [47],[55]). The same is not true, in general, for our incomplete factorization. If one considers the induced splitting splitting need not be convergent. An example is given by the symmetric M-matrix A =B @ For this matrix, the incomplete inverse factorization with a drop tolerance intermediate fill-in is dropped if smaller than T in absolute value) produces an approximate inverse G such that ae(I \Gamma GA) - 1:215 ? 1. This shows that the approximate decomposition cannot be obtained, in general, from an incomplete factorization of A. In this sense, the incomplete inverse factorization is not algebraically equivalent to an incomplete LDU factorization performed on A. 14 Michele Benzi and Miroslav T-uma 5. Fill-in in the biconjugation algorithm. In this section we give a characterization of the fill-in occurring in the factorized inverse obtained by the biconjugation algorithm. These results may serve as a guideline to predict the structure of the factorized approximate inverse, and have an impact on certain aspects of the implementation. It is well-known that structural nonzeros in the inverse matrix A \Gamma1 can be characterized by the paths in the graph of the original matrix A (see [24],[29]). The following lemma states necessary and sufficient conditions for a new entry (fill-in) to be added in one of the z-vectors at the ith step of the biconjugation algorithm. A similar result holds for the w-vectors. We make use of the standard no-cancellation assumption. Lemma 5.1. Let 1 z (i\Gamma1) if and only if l - i, z (i\Gamma1) li and, at the same time, at least one of the two following conditions holds: Proof. Suppose that z (i\Gamma1) 0: Directly from the update formula for the z-vectors we see that z (i\Gamma1) li 6= 0 and l - i, since z (i\Gamma1) lj becomes nonzero in the ith step then clearly p (i\Gamma1) j must be nonzero. But z (i\Gamma1) kj a ik and we get the result. The opposite implication is trivial. 2 Figures 5.1 through 5.6 provide an illustration of the previous lemma. Figure 5.1 shows the nonzero structure of the matrix FS760 1 of order 760 from the Harwell-Boeing collection [21]. Figures 5.2-6 show the structure of the factor Z at different stages of the biconjugation algorithm. These pictures show that in the initial steps, when most of the entries of Z are still zero, the nonzeros in Z are induced by nonzeros in the corresponding positions of A. A similar situation occurs, of course, for the process which computes W . In Figure 5.7 we show the entries of Z which are larger (in absolute in Figure 5.8 we show the incomplete factor - Z obtained with drop tolerance It can be seen how well the incomplete process is able to capture the "large" entries in the complete factor Z. The figures were generated using the routines for plotting sparse matrix patterns from SPARSKIT [50]. Approximate Inverse Preconditioning 15 Figure 5.1-2: Structure of the matrix FS760 1 (left) and of the factor Z (right) after 20 steps of the biconjugation process. Figure 5.3-4: Structure of Z after 70 steps (left) and 200 steps (right) of the biconjugation process. Michele Benzi and Miroslav T-uma Figure 5.5-6: Structure of Z after 400 steps (left) and 760 steps (right) of the biconjugation process. Figure 5.7-8: Structure of entries in Z larger than 10 \Gamma10 (left) and structure of incomplete factor - Z with drop tolerance Approximate Inverse Preconditioning 17 A sufficient condition to have a fill-in in the matrix Z after some steps of the biconju- gation algorithm is given by the following Lemma. Lemma 5.2. Let E) be a bipartite graph with and such that If for some indices i l there is a path in B, then z (i p ) Proof. We use induction on p. Let 0: Of course, z (i 1 \Gamma1) and from Lemma 5.1 we get z (i 1 ) Suppose now that Lemma 5.2 is true for all l ! p. Then, z (i a using the no-cancellation assumption we also have z (i p ) The following theorem gives a necessary and sufficient condition for a nonzero entry to appear in position (l; j), l ! j, in the inverse triangular factor. Theorem 5.3. Let 1 only if for some p - 1 there are Proof. We first show that the stated conditions are sufficient. By Lemma 5.1, the nonzeros a imply that z (i 1 ) l 1 j is also nonzero. If are done. Otherwise, z (i 2 \Gamma1) and a 0: Taking into account that z l 2 we get that z (i 2 ) is nonzero. Repeating these arguments inductively we finally get z (i p ) l Consequently, z (i) Assume now that z lj 6= 0. Lemma 5.1 implies that at least one of the following two conditions holds: either there exists li 0 6= 0, or there exist indices i such that a i 00 k 6= 0, z (i 00 \Gamma1) li 00 6= 0: In the former case we have the necessary conditions. In the latter case we can apply Lemma 5.1 inductively to z (i 00 \Gamma1) After at most j inductive steps we obtain the conditions. 2 Clearly, the characterization of fill-in in the inverse triangular factorization is less transparent than the necessary and sufficient condition which characterize nonzeros in the non- factorized inverse. Michele Benzi and Miroslav T-uma 6. Preconditioning for block triangular matrices. Many sparse matrices arising in real-world applications may be reduced to block triangular form (see Ch. 6 in [20]). In this section we discuss the application of preconditioning techniques to linear systems with a block (lower) triangular coefficient matrix, closely following [30]. The reduction to block triangular form is usually obtained with a two-step procedure, as outlined in [20]. In the first step, the rows of A are permuted to bring nonzero entries on the main diagonal, producing a matrix PA. In the second step, symmetric permutations are used to find the block triangular form [53]. The resulting matrix can be represented as A k1 A k2 \Delta \Delta \Delta A kkC C C C A where the diagonal blocks A ii are assumed to be irreducible. Because A is nonsingular, the diagonal blocks A ii must also be nonsingular. Suppose that we compute approximate inverses of the diagonal blocks A the incomplete biconjugation algorithm, so that A \Gamma1 Z ii ii ii the inverse of A is approximated as follows (cf. [30]): A 22 A k1 A k2 \Delta \Delta QP: The preconditioning step in a conjugate gradient-type method requires the evaluation of the action of G on a vector, i.e. the computation of z = Gd for a given vector d, at each step of the preconditioned iterative method. This can be done by a back-substitution of the z where d =B @ z =B @ Approximate Inverse Preconditioning 19 with the partitioning of - z and - d induced by the block structure of Q(PA)Q T : The computation of which is required by certain preconditioned iterative methods, is accomplished in a similar way. With this approach, fill-in is confined to the approximate inverses of the diagonal blocks, often resulting in a more sparse preconditioner. Notice also that the approximate inverses G ii can be computed in parallel. The price to pay is the loss of part of the explicitness when the approximate inverse preconditioner is applied, as noted in [30]. For comparison purposes, we apply the same scheme with ILU preconditioning. Specif- ically, we approximate A as A 21 A k1 A k2 where each diagonal block A ii is approximated by an ILU decomposition - U ii . Applying the preconditioner requires the solution of a linear system d at each step of the preconditioned iteration. This can be done with the back-substitution where with the same partitioning of - z and - d as above. The use of transposed ILU preconditioning is similar. this type of ILU block preconditioning we introduce some explicitness in the application of the preconditioner. Again, note that the ILU factorizations of the diagonal blocks can be performed in parallel. We will see in the section on numerical experiments that reduction to the block triangular form influences the behavior of the preconditioned iterations in different ways depending on whether approximate inverse techniques or ILU-type preconditioning are used. Michele Benzi and Miroslav T-uma 7. Implementation aspects. It is possible to implement the incomplete inverse factorization algorithm in x4 in at least two distinct ways. The first implementation is similar in spirit to the classical submatrix formulation of sparse Gaussian elimination as represented, for instance, in [19],[57]. This approach relies on sparse incomplete rank-one updates of the matrices - Z and - applied in the form of outer vector products. These updates are the most time-consuming part of the computation. In the course of the updates, new fill-in elements whose magnitude is less than a prescribed drop tolerance T are dropped. In this approach, dynamic data structures have to be used for the - Z and - matrices. Note that at step i of the incomplete inverse factorization, only the ith row a T and the ith column c T are required. The matrix A is stored in static data structures both by rows and by columns (of course, a single array is needed for the numerical values of the entries of A). For this implementation to be efficient, some additional elbow room is necessary. For instance, in the computation of the incomplete - Z factor the elbow room was twice the space anticipated for storing the nonzeros in the factor itself. As we are looking for a preconditioner with about the same number of nonzeros as the original matrix, the estimated number of nonzeros in - Z is half the number of nonzeros in the original matrix A. For each column of - Z we give an initial prediction of fill-in based on the results of x5. Thus, the initial structure of - Z is given by the structure of the upper triangular part of A. Of course, W is handled similarly. If the space initially allocated for a given column is not enough, the situation is solved in a way which is standard when working with dynamic data structures, by looking for a block of free space at the end of the active part of the dynamic data structure large enough to contain the current column, or by a garbage collection (see [57]). Because most of the fill-in in - Z and - W appears in the late steps of the biconjugation process, we were able to keep the amount of dynamic data structure manipulations at relatively low levels. In the following, this implementation will be referred to as the DDS implementation. Despite our efforts to minimize the amount of symbolic manipulations in the DDS im- plementation, some of its disadvantages such as the nonlocal character of the computations and a high proportion of non-floating-point operations still remain. This is an important drawback of submatrix (right-looking, undelayed) algorithms using dynamic data structures when no useful structural prediction is known and no efficient block strategy is used. Even when all the operations are performed in-core, the work with both the row and column lists in each step of the outer cycle is rather irregular. Therefore, for larger problems, most operations are still scattered around the memory and are out-of-cache. As a consequence, it is difficult to achieve high efficiency with the code, and any attempt to parallelize the Approximate Inverse Preconditioning 21 computation of the preconditioner in this form will face serious problems (see [57] for a discussion of the difficulties in parallelizing sparse rank-one updates). For these reasons we considered an alternative implementation (hereafter referred to as SDS) which only makes use of static data structures, based on a left-looking, delayed update version of the biconjugation algorithm. This amounts to a rearrangement of the computations, as shown below. For simplicity we only consider the Z factor, and assume no breakdown occurs: (1) Let z (0)= e 1 z (0) (j \Gamma1) := a T z (j \Gamma1) z (j) := z (j \Gamma1) (j \Gamma1) (j \Gamma1) z (j \Gamma1) := a T z (i\Gamma1) This procedure can be implemented with only static data structures, at the cost of increasing the number of floating-point operations. Indeed, in our implementation we found it necessary to recompute the dot products p (j \Gamma1) z (j \Gamma1) if they are used more than once for updating subsequent columns. This increase in arithmetic complexity is more or less pronounced, depending on the problem and on the density of the preconditioner. On the other hand, this formulation greatly decreases the amount of irregular data structure manipulations. It also appears better suited to parallel implementation, because the dot products and the vector updates in the innermost loop can be done in parallel. Notice that with SDS, it is no longer true that a single row and column of A are used at each step of the outer loop. It is worth mentioning that numerically, the DDS and SDS implementations of the incomplete biconjugation process are completely equivalent. The SDS implementation is straightforward. Suppose the first steps have been completed. In order to determine which columns of the already determined part of - Z play 22 Michele Benzi and Miroslav T-uma a role in the rank-one updates used to form the jth column of - Z we only need a linked list scanning the structure of the columns of A. This linked list is coded similarly to the mechanism which determines the structure of the jth row of the Cholesky factor L in the numerical factorization in SPARSPAK (see [27],[13]). In addition to the approximate inverse preconditioner, we also coded the standard row implementation of the classical ILU(0) preconditioner (see, e.g., [50]). We chose a no-fill implicit preconditioner because we are mostly interested in comparing preconditioners with a nonzero density close to that of the original matrix A. On input, all our codes for the computation of the preconditioners check whether the coefficient matrix has a zero-free diagonal. If not, row reordering of the matrix is used to permute nonzeros on the diagonal. For both the ILU(0) and the approximate inverse factorization, we introduced a simple pivot modification to avoid breakdown. Whenever some diagonal element in any of our algorithms to compute a preconditioner was found to be small, in our case less in absolute value than the IEEE machine precision ffl - 2:2 we increased it to 10 \Gamma3 . We have no special reasons for this choice, other than it worked well in practice. It should be mentioned that in the numerical experiments, this safeguarding measure was required more often for ILU(0) than for the approximate inverse factorization. For the experiments on matrices which can be nontrivially reduced to block triangular form, we used the routine MC13D from MA28 [19] to get the block triangular form. 8. Numerical experiments. In this section we present the results of numerical experiments on a range of problems from the Harwell-Boeing collection [21] and from Tim Davis' collection [17]. All matrices used were rescaled by dividing their elements by the absolute value of their largest nonzero entry. No other scaling was used. The right-hand side of each linear system was computed from the solution vector x of all ones, the choice used, e.g., in [57]. We experimented with several iterative solvers of the conjugate gradient type. Here we present results for three selected methods, which we found to be sufficiently representative: van der Vorst's Bi-CGSTAB method (denoted BST in the tables), the QMR method of Freund and Nachtigal, and Saad and Schultz's GMRES (restarted every 20 steps, denoted G(20) in the tables) with Householder orthogonalization [56]. See [3] for a description of these methods, and the report [9] for experiments with other solvers. Approximate Inverse Preconditioning 23 The matrices used in the experiments come from reservoir simulation (ORS*, PORES2, SAYLR* and SHERMAN*), chemical kinetics (FS5414), network flow (HOR131), circuit simulation (JPWH991, MEMPLUS and ADD*), petroleum engineering (WATT* matrices) and incompressible flow computations (RAEFSKY*, SWANG1). The order N and number NNZ of nonzeros for each test problem are given in Table 1, together with the number of iterations and computing times for the unpreconditioned iterative methods. A y means that convergence was not attained in 1000 iterations for Bi-CGSTAB and QMR, and 500 iterations for GMRES(20). Its Time Table 1: Test problems (N= order of matrix, NNZ= nonzeros in matrix) and convergence results for the iterative methods without preconditioning. Michele Benzi and Miroslav T-uma All tests were performed on a SGI Crimson workstation with RISC processor R4000 using double precision arithmetic. Codes were written in standard Fortran 77 and compiled with the optimization option -O4. CPU time is given in seconds and it was measured using the standard function dtime. The initial guess for the iterative solvers was always x The stopping criterion used was jjr k jj is the (unpreconditioned) updated residual. Note that because r we have that 1 - jjr 0 jj 1 - nzr where nzr denotes the maximum number of nonzeros in a row of A. The following tables present the results of experiments with the ILU(0) preconditioner and with the approximate inverse preconditioner based on the biconjugation process (here- after referred to as AIBC). Observe that the number of nonzeros in the ILU(0) preconditioner is equal to the number NNZ of nonzeros in the original matrix, whereas for the AIBC preconditioner fill-in is given by the total number of nonzeros in the factors - W and - D. In the tables, the number of nonzeros in AIBC is denoted by F ill. Right preconditioning was used for all the experiments. The comparison between the implicit and the explicit preconditioner is based on the amount of fill and on the rate of convergence as measured by the number of iterations. These two parameters can realistically describe the scalar behavior of the preconditioned iterative methods. Of course, an important advantage of the inverse preconditioner, its explicitness, is not captured by this description. The accuracy of the AIBC preconditioner is controlled by the value of the drop tolerance T . Smaller drop tolerances result in a more dense preconditioner and very often (but not always) in a higher convergence rate for the preconditioned iteration. For our experiments we consider relatively sparse preconditioners. In most cases we were able to adjust the value of T so as to obtain an inverse preconditioner with a nonzero density close to that of A (and hence of the ILU(0) preconditioner). Due to the scaling of the matrix entries, the choice very often the right one. We also give results for the approximate inverse obtained with a somewhat smaller value of the drop tolerance, in order to show how the number of iterations can be reduced by allowing more fill-in in the preconditioner. For some problems we could not find a value of T for which the number of nonzeros in AIBC is close to NNZ. In these cases the approximate inverse preconditioner tended to be either very dense or very sparse. Approximate Inverse Preconditioning 25 In Table 2 we give the timings for the preconditioner computation, iteration counts and timings for the three iterative solvers preconditioned with ILU(0). The same information is given in Table 3 for the approximate inverse preconditioner AIBC. For AIBC we give two timings for the construction of the preconditioner, the first for the DDS implementation using dynamic data structures and the second for the SDS implementation using only static data structures. ILU - Its ILU - Time MATRIX P-time BST QMR G(20) BST QMR G(20) RAEFSKY1 2.457 Table 2: Time to form the ILU(0) preconditioner (P-time), number of iterations and time for Bi-CGSTAB, QMR and GMRES(20) with ILU(0) preconditioning. 26 Michele Benzi and Miroslav T-uma P-time AIBC - Its AIBC - Time MATRIX Fill DDS SDS BST QMR G(20) BST QMR G(20) 5204 JPWH991 7063 0.31 0.26 15 27 28 0.24 0.67 0.78 48362 0.68 2.63 33 43 64 2.61 5.39 8.63 26654 0.89 2.45 Table 3: Time to form the AIBC preconditioner (P-time) using DDS and SDS implemen- tations, number of iterations and time for Bi-CGSTAB, QMR and GMRES(20) with AIBC Approximate Inverse Preconditioning 27 It appears from these results that the ILU(0) and AIBC preconditioners are roughly equivalent from the point of view of the rate of convergence, with ILU(0) having a slight edge. On many problems the two preconditioners give similar results. There are a few cases, like PORES2, for which ILU(0) is much better than AIBC, and others (like MEMPLUS) where the situation is reversed. For some problems it is necessary to allow a relatively high fill in the approximate inverse preconditioner in order to have a convergence rate comparable with that insured by ILU(0) (cf. SAYLR4), but there are cases where a very sparse AIBC gives excellent results (see the ADD or the RAEFSKY matrices). It follows that the timings for the iterative part of the solution process are pretty close, on average, for the two preconditioners. We also notice that using a more dense approximate inverse preconditioner (obtained with a smaller value of T ) nearly always reduces the number of iterations, although this does not necessarily mean a reduced computing time since it takes longer to compute the preconditioner and the cost of each iteration is increased. Concerning the matrix PORES2, for which our method gives poor results, we observed that fill-in in the - W factor was very high. We tried to use different drop tolerances for the two factors (the one for - being larger than the one used for - Z) but this did not help. It was observed in [31] that finding a sparse right approximate inverse for PORES2 is very hard and a left approximate inverse should be approximated instead. Unfortunately, our method produces exactly the same approximate inverse (up to transposition) for A and A T , therefore we were not able to cope with this problem effectively. We experienced a similar difficulty with the - W factor for the matrix SHERMAN2. On the other hand, for SHERMAN3 we did not face any of the problems reported in [30] and convergence with the AIBC preconditioner was smooth. As for the time required to compute the preconditioners, it is obvious that ILU(0) can be computed more quickly. On the other hand, the computation of the AIBC preconditioner is not prohibitive. There are problems for which computing AIBC is only two to three times more expensive than computing ILU(0). More important, our experiments with AIBC show that the overall solution time is almost always dominated by the iterative part, unless convergence is extremely rapid, in which case the iteration part takes slightly less time than the computation of the preconditioner. This observation suggests that our approximate inverse preconditioner is much cheaper to construct, in a sequential environment, than approximate inverse preconditioners based 28 Michele Benzi and Miroslav T-uma on the Frobenius norm approach described in x3. Indeed, if we look at the results presented in [30] we see that the sequential time required to construct the preconditioner accounts for a huge portion, often in excess of 90%, of the overall computing time. It is worth emphasizing that the approach based on Frobenius norm minimization and the one we propose seem to produce preconditioners of similar quality, in the sense that they are both comparable with ILU(0) from the point of view of fill-in and rates of convergence, at least on average. As for the different implementations of AIBC, we see from the results in Table 3 that for larger problems, the effect of additional floating-point operations in the SDS implementation is such that the DDS implementation is actually faster. Nevertheless, as already observed the implementation using static data structures may better suited for parallel architec- tures. Because in this paper we only consider a scalar implementation, in the remaining experiments we limit ourselves to the timings for the DDS implementation of AIBC. In all the experiments (excluding the ones performed to measure the timings presented in the tables) we monitored also the "true" residual jjb \Gamma Ax k jj 2 . In general, we found that the discrepancy between this and the norm of the updated residual was small. However, we found that for some very ill-conditioned matrices in the Harwell-Boeing collection (not included in the tables) this difference may be very large. For instance, for some of the LNS* and WEST* matrices, we found that jjr k jj for the final value of r k . This happened both with the ILU(0) and with the approximate inverse preconditioner, and we regarded this as a failure of the preconditioned iterative method. We present in Tables 4 and 5 the results of some experiments on matrices which have been reduced to block lower triangular form. We compared the number of iterations of the preconditioned iterative methods and their timings for the block approximate inverse preconditioner and for the block ILU(0) preconditioner as described in x6. Since some of the matrices have only trivial block lower triangular form (one block, or two blocks with one of the blocks of dimension one for some matrices) we excluded them from our experiments. In Table 4 we give for each matrix the number NBL of blocks and the results of experiments with ILU(0). In Table 5 we give analogous results for the AIBC preconditioner. The amount of fill-in (denoted by F ill) for AIBC is computed as the fill-in in the approximate inverses of the diagonal blocks plus the number of nonzero entries in the off-diagonal blocks. Approximate Inverse Preconditioning 29 Block ILU - Its Block ILU - Time Table 4: Time to compute the block ILU preconditioner (P-time), number of iterations and time for Bi-CGSTAB, QMR and GMRES(20) with block ILU(0) preconditioning. Block AIBC - Its Block AIBC - Time MATRIX Fill P-time BST QMR G(20) BST QMR G(20) Table 5: Time to compute the block AIBC preconditioner (P-time) , number of iterations and time in seconds for Bi-CGSTAB, QMR and GMRES(20) with block AIBC preconditioning. It is clear that in general the reduction to block triangular form does not lead to a noticeable improvement in the timings, at least in a sequential implementation. We observe that when the block form is used, the results for ILU(0) are sometimes worse. This can Michele Benzi and Miroslav T-uma probably be attributed to the permutations, which are known to cause in some cases a degradation of the rate of convergence of the preconditioned iterative method [22]. A notable exception is the matrix WATT2, for which the number of iterations is greatly reduced. On the other hand, the results for the block approximate inverse preconditioner are mostly unchanged or somewhat better. Again, matrix WATT2 represents an exception: this problem greatly benefits from the reduction to block triangular form. In any case, permutations did not adversely affect the rate of convergence of the preconditioned iterative method. This fact suggests that perhaps the approximate inverse preconditioner is more robust than ILU(0) with respect to reorderings. To gain more insight on how permutations of the original matrix can influence the quality of both types of preconditioners, we did some experiments where the matrix A was permuted using the minimum degree algorithm on the structure of A + A T (see [28]). We applied the resulting permutation to A symmetrically to get PAP T , in order to preserve the nonzero diagonal. Tables 6 and 7 present the results for the test matrices having trivial block triangular form. The corresponding preconditioners are denoted by ILU(0)-MD and AIBC-MD, respectively. ILU-MD - Its ILU-MD - Time MATRIX P-time BST QMR G(20) BST QMR G(20) 26 43 47 2.18 4.57 6.45 Table Time to compute the ILU(0) preconditioner (P-time) for A permuted according to minimum degree algorithm on A number of iterations and time for Bi-CGSTAB, QMR and GMRES(20) with ILU(0)-MD preconditioning. Approximate Inverse Preconditioning 31 AIBC-MD - Its AIBC-MD - Time MATRIX Fill P-time BST QMR G(20) BST QMR G(20) 7152 0.31 43 48 0.38 1.23 1.38 19409 0.58 104 95 y 2.91 4.14 y Table 7: Time to compute the AIBC preconditioner (P-time) for A permuted by the minimum degree algorithm on A number of iterations and time for Bi-CGSTAB, QMR and GMRES(20) with AIBC-MD preconditioning. The results in Table 6 show that for some problems, especially those coming from PDEs, minimum degree reordering has a detrimental effect on the convergence of the iterative solvers preconditioned with ILU(0). In some cases we see a dramatic increase in the number of iterations. This is in analogy with the observed fact (see, e.g., [22]) that when the minimum degree ordering is used, the no-fill incomplete Cholesky decomposition of an SPD Michele Benzi and Miroslav T-uma matrix is a poor approximation of the coefficient matrix, at least for problems arising from the discretization of 2D PDEs. The convergence of the conjugate gradient method with such a preconditioner (ICCG(0)) is much slower than if the natural ordering of the unknowns was used. Here we observe a similar phenomenon for nonsymmetric linear systems. Note the rather striking behavior of matrix ADD20, which benefits greatly from the minimum degree reordering (this matrix arises from a circuit model and not from the discretization of a PDE). It was also observed in [22] that the negative impact of minimum degree on the rate of convergence of PCG all but disappears when the incomplete Cholesky factorization of A is computed by means of a drop tolerance rather than by position. It is natural to ask whether the same holds true for the approximate inverse preconditioner AIBC, which is computed using a drop tolerance. The results in Table 7 show that this is indeed the case. For most of the test problems the number of iterations was nearly unaffected (or better) and in addition we note that the minimum degree ordering helps in preserving sparsity in the incomplete inverse factors. While this is usually not enough to decrease the computing times, the fact that it is possible to reduce storage demands for the approximate inverse preconditioner without negatively affecting the convergence rates might become important for very large problems. We conclude this section with some observations concerning the choice of the drop tolerance T . In all our experiments we used a fixed value of T throughout the incomplete biconjugation process. However, relative drop tolerances, whose value is adapted from step to step, could also be considered (see [57] for a thorough discussion of the issues related to the choice of drop tolerances in the context of ILU). We have observed that the amount of fill-in is distributed rather unevenly in the course of the approximate inverse factorization. A large proportion of nonzeros is usually concentrated in the last several columns of - Z and - W . For some problems with large fill, it may be preferable to switch to a larger drop tolerance when the columns of the incomplete factors start filling-in strongly. Conversely, suppose we have computed an approximate inverse preconditioner for a certain value of T , and we find that the preconditioned iteration is converging slowly. Provided that enough storage is available, one could then try to recompute at least some of the columns of - Z and - using a smaller value of T . Unfortunately, for general sparse matrices there is no guarantee that this will result in a preconditioner of improved quality. Indeed, allowing more nonzeros in the preconditioner does not always result in a reduced number of iterations. Approximate Inverse Preconditioning 33 Finally, it is worthwhile to observe that a dual threshold variant of the incomplete inverse factorization could be adopted, see [51]. In this approach, a drop tolerance is applied but a maximum number of nonzeros per column is specified and enforced during the computation of the preconditioner. In this way, it is possible to control the maximum storage needed by the preconditioner, which is important for an automated implementation. This approach has not been tried yet, but we hope to do so in the near future. 9. Conclusions and future work. In this paper we have developed a sparse approximate inverse preconditioning technique for nonsymmetric linear systems. Our approach is based on a procedure to compute two sets of biconjugate vectors, performed incompletely to preserve sparsity. This algorithm produces an approximate triangular factorization of A \Gamma1 , which is guaranteed to exist if A is an H-matrix (similar to the ILU factorization). The factorized sparse approximate inverse is used as an explicit preconditioner for conjugate gradient-type methods. Applying the preconditioner only requires sparse matrix-vector products, which is of considerable interest for use on parallel computers. The new preconditioner was used to enhance the convergence of different iterative solvers. Based on extensive numerical experiments, we found that our preconditioner can insure convergence rates which are comparable, on average, with those from the standard preconditioner. While the approximate inverse factorization is more time-consuming to compute than ILU(0), its cost is not prohibitive, and is typically dominated by the time required by the iterative part. This is in contrast with other approximate inverse preconditioners, based on Frobenius norm minimization, which produce similar convergence rates but are very expensive to compute. It is possible that in a parallel environment the situation will be reversed, since the preconditioner construction with the Frobenius norm approach is inherently parallel. How- ever, there is some scope for parallelization also in the inverse factorization on which our method is based: for instance, the approximate inverse factors - Z and - W can be computed largely independent of each other. Clearly, this is a point which requires further research, and no conclusion can be drawn until parallel versions of this and other approximate inverse preconditioners have been implemented and tested. Our results point to the fact that the quality of the approximate inverse preconditioner is not greatly affected by reorderings of the coefficient matrix. This is important in practice because it suggests that we may use permutations to increase the potential for parallelism or to reduce the amount of fill in the preconditioner, without spoiling the rate of convergence. 34 Michele Benzi and Miroslav T-uma The theoretical results on fill-in in x5 provide guidelines for the use of pivoting strategies for enhancing the sparsity of the approximate inverse factors, and this is a topic that deserves further research. Based on the results of our experiments, we conclude that the technique introduced in this paper has the potential to become a useful tool for the solution of large sparse nonsymmetric linear systems on modern high-performance architectures. Work on a parallel implementation of the new preconditioner is currently under way. Future work will also include a dual threshold implementation of the preconditioner computation. Acknowledgments . We would like to thank one of the referees for helpful comments and suggestions, and Professor Miroslav Fiedler for providing reference [24]. The first author gratefully acknowledges the hospitality and excellent research environment provided by the Institute of Computer Science of the Czech Academy of Sciences. --R Parallel Implementation of Preconditioned Conjugate Gradient Methods for Solving Sparse Systems of Linear Equations. Iterative Solution Methods. Templates for the Solution of Linear Systems. Parallel algorithms for the solution of certain large sparse linear systems. A Direct Row-Projection Method for Sparse Linear Systems A direct projection method for sparse linear systems. An explicit preconditioner for the conjugate gradient method. A sparse approximate inverse preconditioner for the conjugate gradient method. A sparse approximate inverse preconditioner for nonsymmetric linear systems. Krylov methods preconditioned with incompletely factored matrices on the CM-2 Approximate inverse preconditioners for general sparse matrices. Approximate inverse techniques for block-partitioned matrices User's guide for SPARSPAK-A: Waterloo sparse linear equations package Block preconditioning for the conjugate gradient method. Approximate inverse preconditionings for sparse linear systems. Sparse matrix collection. Decay rates for inverses of band matrices. Direct Methods for Sparse Matrices. Users' guide for the Harwell-Boeing sparse matrix collection The effect of ordering on preconditioned conjugate gradients. A stability analysis of incomplete LU factorizations. Inversion of bigraphs and connection with the Gauss elimination. An Introduction to Numerical Linear Algebra. Notes on the solution of algebraic linear simultaneous equations. Computer Solution of Large Sparse Positive Definite Systems. The evolution of the minimum degree algorithm. Predicting structure in sparse matrix computations. On approximate-inverse preconditioners Parallel preconditioning with sparse approximate inverses. Parallel preconditioning and approximate inverses on the Connection Machine. A parallel preconditioned conjugate gradient package for solving sparse linear systems on a Cray Y-MP Inversion of matrices by biorthogonalization and related results. Method of conjugate gradients for solving linear systems. The Theory of Matrices in Numerical Analysis. Polynomial preconditioning for conjugate gradient calculations. The efficient parallel iterative solution of large sparse linear sys- tems Explicitly preconditioned conjugate gradient method for the solution of unsymmetric linear systems. New convergence results and preconditioning strategies for the conjugate gradient method. Factorized sparse approximate inverse (FSAI) preconditionings for solving 3D FE systems on massively parallel computers II: Iterative construction of FSAI preconditioners. Factorized sparse approximate inverse preconditioning I: Theory. Factorized sparse approximate inverse preconditioning II: Solution of 3D FE systems on massively parallel computers. Krylov Methods for the Numerical Solution of Initial-Value Problems in Differential-Algebraic Equations An incomplete factorization technique for positive definite linear systems. An iterative solution method for linear systems of which the coefficient matrix is a symmetric M-matrix Some properties of approximate inverses of matrices. Preconditioning techniques for nonsymmetric and indefinite linear systems. SPARSKIT: A basic tool kit for sparse matrix computations. ILUT: A dual threshold incomplete LU factorization. Conjugate direction methods for solving systems of linear equations. High performance preconditioning. Matrix Iterative Analysis. Implementation of the GMRES method using Householder transformations. Computational Methods for General Sparse Matrices. --TR --CTR Kai Wang , Jun Zhang, Multigrid treatment and robustness enhancement for factored sparse approximate inverse preconditioning, Applied Numerical Mathematics, v.43 n.4, p.483-500, December 2002 Claus Koschinski, New methods for adapting and for approximating inverses as preconditioners, Applied Numerical Mathematics, v.41 n.1, p.179-218, April 2002 Stephen T. Barnard , Luis M. Bernardo , Horst D. Simon, An MPI Implementation of the SPAI Preconditioner on the T3E, International Journal of High Performance Computing Applications, v.13 n.2, p.107-123, May 1999 N. Guessous , O. Souhar, Multilevel block ILU preconditioner for sparse nonsymmetric M-matrices, Journal of Computational and Applied Mathematics, v.162 n.1, p.231-246, 1 January 2004 Matthias Bollhfer , Volker Mehrmann, Some convergence estimates for algebraic multilevel preconditioners, Contemporary mathematics: theory and applications, American Mathematical Society, Boston, MA, 2001 Michele Benzi , Miroslav Tma, A parallel solver for large-scale Markov chains, Applied Numerical Mathematics, v.41 n.1, p.135-153, April 2002 Mansoor Rezghi , S. Mohammad Hosseini, An ILU preconditioner for nonsymmetric positive definite matrices by using the conjugate Gram-Schmidt process, Journal of Computational and Applied Mathematics, v.188 n.1, p.150-164, 1 April 2006 M. H. Koulaei , F. Toutounian, On computing of block ILU preconditioner for block tridiagonal systems, Journal of Computational and Applied Mathematics, v.202 n.2, p.248-257, May, 2007 Michele Benzi, Preconditioning techniques for large linear systems: a survey, Journal of Computational Physics, v.182 n.2, p.418-477, November 2002

approximate inverses;incomplete factorizations;sparse linear systems;conjugate gradient-type methods;sparse matrices;preconditioning

292377

Approximate Inverse Preconditioners via Sparse-Sparse Iterations.

The standard incomplete LU (ILU) preconditioners often fail for general sparse indefinite matrices because they give rise to "unstable" factors L and U. In such cases, it may be attractive to approximate the inverse of the matrix directly. This paper focuses on approximate inverse preconditioners based on minimizing ||I-AM||F, where AM is the preconditioned matrix. An iterative descent-type method is used to approximate each column of the inverse. For this approach to be efficient, the iteration must be done in sparse mode, i.e., with "sparse-matrix by sparse-vector" operations. Numerical dropping is applied to maintain sparsity; compared to previous methods, this is a natural way to determine the sparsity pattern of the approximate inverse. This paper describes Newton, "global," and column-oriented algorithms, and discusses options for initial guesses, self-preconditioning, and dropping strategies. Some limited theoretical results on the properties and convergence of approximate inverses are derived. Numerical tests on problems from the Harwell--Boeing collection and the FIDAP fluid dynamics analysis package show the strengths and limitations of approximate inverses. Finally, some ideas and experiments with practical variations and applications are presented.

Introduction . The incomplete LU factorization preconditioners were originally developed for M-matrices that arise from the discretization of very simple partial differential equations of elliptic type, usually in one variable. For the rather common situation where the matrix A is indefinite, standard ILU factorizations may face several difficulties, the best known of which is the encounter of a zero pivot. However, there are other problems that are just as serious. Consider an incomplete factorization of the form where E is the error. The preconditioned matrices associated with the different forms of preconditioning are similar to What is sometimes missed is the fact that the error matrix E in (1.1) is not as important as the preconditioned error matrix L shown in (1.2) above. When the matrix A is diagonally dominant, L and U are typically well conditioned, and the size of L remains confined within reasonable limits, typically with a clustering of its eigenvalues around the origin. On the other hand, when the original matrix is not diagonally dominant, L \Gamma1 or U \Gamma1 may have very large norms, causing the error to be very large and thus adding large perturbations to the identity matrix. This form of instability was studied by Elman [14] in a detailed analysis of ILU and MILU preconditioners for finite difference matrices. It can be observed experimentally This work was supported in part by the National Science Foundation under grant NSF/CCR- 9214116 and in part by NASA under grant NAG2-904. y Department of Computer Science and Minnesota Supercomputer Institute, University of Minnesota, 4-192 EE/CSci Bldg., 200 Union St., S.E., Minneapolis, Minnesota, 55455-0154 ([email protected] and [email protected]). that ILU preconditioners can be very poor when L \Gamma1 or U \Gamma1 are large, and that this situation often occurs for indefinite problems, or problems with large nonsymmetric parts. One possible remedy that has been proposed is stabilized or perturbed incomplete factorizations, for example [15] and the references in [25]. A numerical comparison with these preconditioners will be given later. In this paper, we consider trying to find a preconditioner that does not require solving a linear system. For example, we can precondition the original system with a sparse matrix M that is a direct approximation to the inverse of A. Sparse approximate inverses are also necessary for incomplete block factorizations with large sparse blocks, as well as several other applications, also described later. We focus on methods of finding approximate inverses based on minimizing the Frobenius norm of the residual matrix I \Gamma AM , first suggested by Benson and Frederickson [5, 6]. Consider the minimization of to seek a right approximate inverse. An important feature of this objective function is that it can be decoupled as the sum of the squares of the 2-norms of the individual columns of the residual matrix I \Gamma AM in which e j and m j are the j-th columns of the identity matrix and of the matrix M , respectively. Thus, minimizing (1.4) is equivalent to minimizing the individual functions This is clearly useful for parallel implementations. It also gives rise to a number of different options. The minimization in (1.5) is most often performed directly by prescribing a sparsity pattern for M and solving the resulting least squares problems. Grote and Simon choose M to be a banded matrix with 2p emphasizing the importance of the fast application of the preconditioner in a CM-2 implementa- tion. This choice of structure is particularly suitable for banded matrices. Cosgrove, D'iaz and Griewank [10] select the initial structure of M to be diagonal and then use a procedure to improve the minimum by updating the sparsity pattern of M . New fill-in elements are chosen so that the fill-in contributes a certain improvement while minimizing the number of new rows in the least squares subproblem. In similar work by Grote and Huckle [18], the reduction in the residual norm is tested for each candidate fill-in element, but fill-in may be introduced more than one at a time. In other related work, Kolotilina and Yeremin [23] consider symmetric, positive definite systems and construct factorized sparse approximate inverse preconditioners which are also symmetric, positive definite. Each factor implicitly approximates the inverse of the lower triangular Cholesky factor of A. The structure of each factor is chosen to be the same as the structure of the lower triangular part of A. In their more recent work [24], fill-in elements may be added, and their locations are chosen such that the construction and application of the approximate inverse is not much more expensive on a model hypercube computer. Preconditioners for general systems may be constructed by approximating the left and right factors separately. This paper is organized as follows. In x2, we present several approximate inverse algorithms based on iterative procedures, as well as describe sparse-sparse implementation and various options. We derive some simple theoretical results for approximate inverses and the convergence behavior of the algorithms in x3. In x4, we show the strengths and limitations of approximate inverse preconditioners through numerical tests with problems from the Harwell-Boeing collection and the FIDAP fluid dynamics analysis package. Finally in x5, we present some ideas and experiments with practical variations and applications of approximate inverses. 2. Construction of the approximate inverse via iteration. The sparsity pattern of an approximate inverse of a general matrix should not be prescribed, since an appropriate pattern is usually not known beforehand. In contrast to the previous work described above, the locations and values of the nonzero elements are determined naturally as a side-effect of utilizing an iterative procedure to minimize (1.3) or (1.5). In addition, elements in the approximate inverse may be removed by a numerical dropping strategy if they contribute little to the inverse. These features are clearly necessary for general sparse matrices. In xx2.1 and 2.2 we briefly describe two approaches where M is treated as a matrix in its entirety, rather than as individual columns. We found, however, that these methods converge more slowly than if the columns are treated separately. In the remaining sections, we consider this latter approach and the various options that are available. 2.1. Newton iteration. As an alternative to directly minimizing the objective function (1.3), an approximate inverse may also be computed using an iterative process known as the method of Hotelling and Bodewig [20]. This method, which is modeled after Newton's method for solving f(x) j 1=x \Gamma a = 0, has many similarities to our descent methods which we describe later. The iteration takes the form For convergence, we require that the spectral radius of I \Gamma AM 0 be less than one, and if we choose an initial guess of the form M convergence is achieved if In practice, we can follow Pan and Reif [27] and use for the right approximate inverse. As the iterations progress, M becomes denser and denser, and a natural idea here is to perform the above iteration in sparse mode [26], i.e., drop some elements in M or else the iterations become too expensive. In this case, however, the convergence properties of the Newton iteration are lost. We will show the results of some numerical experiments in x4. 2.2. Global iteration. In this section we describe a 'global' approach to minimizing (1.3), where we use a descent-type method, treating M as an unknown sparse matrix. The objective function (1.3) is a quadratic function on the space of n \Theta n matrices, viewed as objects in R n 2 . The actual inner product on the space of matrices with which the function (1.4) is associated is One possible descent-type method we may use is steepest descent which we will describe later. In the following, we will call the array representation of an n 2 vector X the n \Theta n matrix whose column vectors are the successive n-vectors of X . In descent algorithms a new iterate M new is defined by taking a step along a selected direction G, i.e., in which ff is selected to minimize the objective function associated with M new . This is achieved by taking hAG; AGi tr ((AG) T AG) is the residual matrix. Note that the denominator may be computed as kAGk 2 F . After each of these descent steps is taken, the resulting matrix will tend to become denser. It is therefore essential to apply some kind of numerical dropping, either to the new M or to the search direction G before taking the descent step. In the first case, the descent nature of the step is lost, i.e., it is no longer guaranteed that F (Mnew ) F (M ), while in the second case, the fill-in in M is more difficult to control. We will discuss both these alternatives in x2.5. The simplest choice for the descent direction G is to take it to be the residual is the new iterate. The corresponding descent algorithm is referred to as the Minimal Residual (MR) algorithm. In the simpler case where numerical dropping is applied to M , our global Minimal Residual algorithm will have the following form. Algorithm 2.1. (Global Minimal Residual descent algorithm) 1. Select an initial M 2. Until convergence do 3. Compute G 4. Compute ff by (2.2) 5. Compute 6. Apply numerical dropping to M 7. End do Another popular choice is to take G to be the direction of steepest descent, i.e., the direction opposite to the gradient. Thinking in terms of n 2 vectors, the gradient of F can be viewed as an n 2 vector g such that where (\Delta; \Delta) is the usual Euclidean inner product. If we represent all vectors as 2- dimensional n \Theta n arrays, then the above relation is equivalent to This allows us to determine the gradient as an operator on arrays, rather than n 2 vectors, as is done in the next proposition. Proposition 2.2. The array representation of the gradient of F with respect to M is the matrix in which R is the residual matrix Proof. For any matrix E we have \Theta (R \Gamma AE) T (R \Theta (AE) T R +R Thus, the differential of F applied to E is the inner product of \Gamma2A T R with E plus a second order term. The gradient is therefore simply \Gamma2A T R. The steepest descent algorithm consists of simply replacing G in line 3 of the MR algorithm described above by This algorithm can be a very slow in some cases, since it is essentially a steepest descent-type algorithm applied to the normal equations. In either global steepest descent or minimal residual, we need to form and store the G matrix explicitly. The scalars kAGk 2 F and tr(R T AG) can be computed from the successive columns of AG, which can be generated, used, and discarded. Therefore, we need not store the matrix AG. We will show the results of some numerical experiments with this global iteration and compare them with other methods in x4. 2.3. Implementation of sparse mode MR and GMRES. We now describe column-oriented algorithms which consist of minimizing the individual objective functions (1.5). We perform this minimization by taking a sparse initial guess and solving approximately the n linear subproblems with a few steps of a nonsymmetric descent-type method, such as MR or untruncated GMRES. For this method to be efficient, the iterative method must work in sparse mode, i.e., m j is stored and operated on as a sparse vector, and the Arnoldi basis in GMRES is kept in sparse format. In the following MR algorithm, n i iterations are used to solve (2.3) approximately for each column, giving an approximation to the j-th column of the inverse of A. Each initial m j is taken from the columns of an initial guess, M 0 . Again, we assume numerical dropping is applied to M . In the GMRES version of the algorithm, we never use restarting since since n i is typically very small. Also, a variant called FGMRES [31] which allows an arbitrary Arnoldi basis, is actually used in this case. Algorithm 2.3. (Minimal Residual iteration) 1. 2. For each column do 3. 4. For do 6 E. CHOW AND Y. SAAD 5. r j := e 7. 8. Apply numerical dropping to m j 9. End do 10. End do Thus, the algorithm computes the current residual r j and then minimizes the residual norm e j;new in the set In the sparse implementation of MR and GMRES, the matrix-vector product, SAXPY, and dot product kernels now all entirely involve sparse vectors. The matrix-vector product is much more efficient if the sparse matrix is stored by columns since all the entries do not need to be traversed. Efficient codes for all these kernels may be constructed which utilize a full n-length work vector [11]. Columns from an initial guess M 0 for the approximate inverse are used as the initial guesses for the iterative solution of the linear subproblems. There are two obvious choices: M . The scale factor ff is chosen to minimize the spectral radius ae(I \Gamma ffAM ). Denoting the initial guess as M writing @ @ff leads to tr(AM) The transpose initial guess is more expensive to use because it is denser than the identity initial guess. However, for very indefinite systems, this guess immediately produces a symmetric positive definite preconditioned system, corresponding to the normal error equations. Depending on the structure of the inverse, a denser initial guess is often required to involve more of the matrix A in the computation. Interest- ingly, the cheaper the computation, the more it uses only 'local' information, and the less able it may be to produce a good approximate inverse. The choice of initial guess also depends to some degree on 'self-preconditioning' which we describe next. Additional comments on the choice of initial guess will be presented there. 2.4. Self-preconditioning. The approximate solution of the linear subproblems using an iterative method suffers from the same problems as solving the original problem if A is indefinite or poorly conditioned. However, the linear systems may be preconditioned with the columns that have already been computed. More precisely, each system (2.3) for approximating column j may be preconditioned with 0 where the first 0 are the m k that already have been computed, and the remaining columns are the initial guesses for the m k , j k n. This suggests that it is possible to define outer iterations that sweep over the matrix, as well as inner iterations that compute each column. On each subsequent outer iteration, the initial guess for each column is the previous result for that column. This technique usually results in much faster convergence of the approximate inverse. Unfortunately with this approach, the parallelism of constructing the columns of the approximate inverse simultaneously is lost. However, there is another variant of self-preconditioning that is easier to implement and more easily parallelizable. Quite simply, all the inner iterations are computed simultaneously and the results of all the columns are used as the self-preconditioner for the next outer iteration. Thus, the preconditioner for the inner iterations changes only after each outer iteration. The performance of this variant usually lies between full self-preconditioning and no self- preconditioning. A more reasonable compromise is to compute blocks of columns in parallel, and some (inner) self-preconditioning may be used. Self-preconditioning is particularly valuable for very indefinite problems when combined with a scaled transpose initial guess; the initial preconditioned system AM 0 is positive definite, and the subsequent preconditioned systems somewhat maintain this property, even in the presence of numerical dropping. Self-preconditioning with a transpose initial guess, however, may produce worse results if the matrix A is very ill-conditioned. In this case, the initial worsening of the conditioning of the system is too severe, and the alternative scaled identity initial guess should be used instead. We have also found cases where self-preconditioning produces worse results, usually for positive definite problems; this is not surprising, since the minimizations would progress very well, only to be hindered by self-preconditioning with a poor approximate inverse in the early stages. Numerical evidence of these phenomena will be provided in x4. Algorithm 2.4 implements the Minimal Residual iteration with self-preconditioning. In the algorithm, n iterations and n i inner iterations are used. Again, initially. We have also indicated where numerical dropping might be applied. Algorithm 2.4. (Self-preconditioned Minimal Residual iteration) 1. Start: 2. For 3. For each column do 4. Define s := 5. For do As 7. z := Mr 8. q := Az 9. ff := (r;q) 11. Apply numerical dropping to s 12. End do 13. Update j-th column of 14. End do 15. End do In a FORTRAN 77 implementation, M is stored as n sparse vectors, each holding up to lfil entries. M is thus constructed in place. The multiple outer iterations used in constructing the approximate inverse suggests the use of factorized updates. Factorized matrices can express denser matrices than the sum of their numbers of elements alone. Suppose that one outer iteration has produced the approximate inverse M 1 . Then a second outer iteration tries to find M 2 , an approximate inverse to AM 1 . In general, after i outer iterations, we are looking for the update M i+1 which minimizes min It is also possible to construct factorized approximate inverses of the form min F which alternate from left to right factors. This latter form is reminiscent of the symmetric form of Kolotilina and Yeremin [23]. Since the product never formed explicitly, the factorized approach effectively uses less memory for the preconditioner at the cost of multiplying with each factor for each matrix-vector multiplication. This approach may be suitable for very large problems, where memory rather than solution time is the limiting factor. The implementation, however, is much more complex, since a sequence of matrices needs to be maintained. 2.5. Numerical dropping strategies. There are many options for numerical dropping. So far, to ease the presentation, we have only discussed the case where dropping is performed on the solution vectors or matrices. Section 2.5.1 discusses this case in more detail, while x2.5.2 discusses the case where dropping is applied to the search directions. In the latter case, the descent property of the algorithms is maintained. 2.5.1. Dropping in the solution. When dropping is performed on the solution, we have options for 1. when dropping is performed, and 2. which elements are dropped. In the previous algorithms, we have made the first point precise; however, there are other alternatives. For example, dropping may be performed only after M or each column of M is computed. Typically this option is too expensive, but as a compromise, dropping may be performed at the end of a few inner iterations, before M is updated, namely before step 13 in Algorithm 2.4. Interestingly, we found experimentally that this option is not always better. In GMRES, the Krylov basis vectors are kept sparse by dropping elements just after the self-preconditioning step, before the multiplication by A. To address which elements are dropped, we can utilize a dual threshold strategy based on a drop tolerance, droptol, and the maximum number of elements per column, lfil. By limiting the maximum number of elements per column, the maximum storage for the preconditioner is known beforehand. The drop tolerance may be applied directly to the elements to be dropped: i.e., elements are dropped if their magnitude is smaller than droptol. However, we found that this strategy could cause spoiling of the minimization, i.e., the residual norm may increase after several steps, along with a deterioration of the quality of the preconditioner. dropping small elements in m j is sub-optimal, one may ask the question whether or not dropping can be performed more optimally. A simple perturbation analysis will help understand the issues. We denote by m j the current column, and by " the perturbed column formed by adding the sparse column d in the process of numerical dropping. The new column and corresponding residual are therefore The square of the residual norm of the perturbed m j is given by Recall that \Gamma2A T r j is the gradient of the function (1.5). As is expected from standard results in optimization, if d is in the direction opposite to the gradient, and if it is small enough, we can achieve a decrease of the residual norm. Spoiling occurs when close to zero so that for practical sizes of kdk 2 , kAdk 2 becomes dominant, causing an increase in the residual norm. Consider specifically the situation where only one element is dropped, and assume that all the columns Ae i of A have been pre-scaled so that kAe i 1. In this case, and the above equation becomes A strategy could therefore be based on attempting to make the function nonpositive, a condition which is easy to verify. This suggests selecting elements to drop in m j only at indices i where the selection function (2.8) is zero or negative. However, note that this is not entirely rigorous since in practice a few elements are dropped at the same time. Thus we do not entirely perform dropping via numerical values alone. In a two-stage process, we first select a number of candidate elements to be dropped based only on the numerical size as determined by a certain tolerance. Among these, we drop all those that satisfy the condition or we can keep those lfil elements that have the largest ae ij . Another alternative is based on attempting to achieve maximum reduction in the function (2.8). Ideally, we wish to have since this will achieve the 'optimal' reduction in (2.8) This leads to the alternative strategy of dropping elements in positions i of m j where are the smallest. We found, however, that this strategy produces poorer results than the previous one, and neither of these strategies completely eliminate spoiling. 2.5.2. Dropping in the search direction. Dropping may be performed on the search direction G in Algorithm 2.1, or equivalently in r j and z in Algorithms 2.3 and 2.4 respectively. In these cases, the descent property of the algorithms is maintained, and the problem of spoiling is avoided. Starting with a sparse initial guess, the allowed number of fill-ins is gradually increased at each iteration. For an MR-like algorithm, the search direction d is derived by dropping entries from the residual direction r. So that the sparsity pattern of the solution x is controlled, d is chosen to have the same sparsity pattern as x, plus one new entry, the largest entry in absolute value. No drop tolerance is used. Minimization is performed by choosing the step-length as (Ad; Ad) and thus the residual norm for the new solution is guaranteed to be not more than the previous residual norm. In contrast to Algorithm 2.3, the residual may be updated with very little cost. The iterations may continue as long as the residual norm is larger than some threshold, or a set number of iterations may be used. If A is indefinite, the normal equations residual direction A T r may be used as the search direction, or simply to determine the location of the new fill-in. It is interesting to note that the largest entry in A T r gives the greatest residual norm reduction in a one-dimensional minimization. When fill-in is allowed to increase gradually using this search direction, this technique becomes very similar to the adaptive selection scheme of [18]. The effect is also similar to self-preconditioning with a transpose initial guess. At the end of each iteration, it is possible to use a second stage that exchanges entries in the solution with new entries if this causes a reduction in the residual norm. This is required if the sparsity pattern in the approximate inverse needs to change as the approximations progress. We have found this to be necessary, particularly for very unstructured matrices, but have not yet found a strategy that is genuinely effective [7]. As a result, approximations using numerical dropping in the solution are often better, even though the scheme just described has a stronger theoretical justification, similar to that of [18]. This also shows that the adaptive scheme of [18] may benefit from such an exchange strategy. Algorithm 2.5 implements a Minimal Residual-like algorithm with this numerical dropping strategy. The number of inner iterations is usually chosen to be lfil or somewhat larger. Algorithm 2.5. (Self-preconditioned MR algorithm with dropping in search direction) 1. Start: 2. For each column do 3. 4. r j := e 5. For do 7. Choose d to be t with the same pattern as If one entry which is the largest remaining entry in absolute value 8. q := Ad 9. ff := (r j ;q) 12. End do 13. End do If dropping is applied to the unpreconditioned residual, then economical use of this approximate inverse technique is not limited to approximating the solution to linear systems with sparse coefficient matrices or sparse right-hand sides. An approximation may be found, for example, to a factorized matrix, or a dense operator which may only be accessed with a matrix-vector product. Such a need may arise, for instance, when preconditioning row projection systems. These approximations are not possible with other existing approximate inverse techniques. We must mention here that any adaptive strategy such as this one for choosing the sparsity pattern makes massive parallelization of the algorithm more difficult. If, for instance, each processor has the task of computing a few columns of the approximate inverse, it is not known beforehand which columns of A must be fetched into each processor. 2.6. Cost of constructing the approximate inverse. The cost of computing the approximate inverse is relatively high. Let n be the dimension of the linear system, be the number of outer iterations, and n i be the number of inner iterations (n in Algorithm 2.5). We approximate the cost by the number of sparse matrix-sparse vector multiplications in the sparse mode implementation of MR and GMRES. Profiling for a few problems shows that this operation accounts for about three-quarters of the time when self-preconditioning is used. The remaining time is used primarily by the sparse dot product and sparse SAXPY operations, and in the case of sparse mode GMRES, the additional work within this algorithm. If Algorithm 2.4 is used, two sparse mode matrix-vector products are used, the first one for computing the residual; three are required if self-preconditioning is used. In Algorithm 2.5 the residual may be updated easily and stored, or recomputed as in Algorithm 2.4. Again, an additional product is required for self-preconditioning. The cost is simply nn times the number of these sparse mode matrix-vector multipli- cations. Each multiplication is cheap, depending on the sparseness of the columns in M . Dropping in the search directions, however, is slightly more expensive because, although the vectors are sparser at the beginning, it typically requires much more inner iterations (e.g., one for each fill-in). In Newton iteration, two sparse matrix-sparse matrix products are required, although the convergence rate may be doubled with form of Chebyshev acceleration [28]. Global iterations without self-preconditioning require three matrix-matrix products. These costs are comparable to the column-oriented algorithms above. 3. Theoretical considerations. Theoretical results regarding the quality of approximate inverse preconditioners are difficult to establish. However, we can prove a few rather simple results for general approximate inverses and the convergence behavior of the algorithms. 3.1. Nonsingularity of M . An important question we wish to address is whether or not an approximate inverse obtained by the approximations described earlier can be singular. It cannot be proved that M is nonsingular unless the approximation is accurate enough, typically to a level that is impractical to attain. This is a difficulty for all approximate inverse preconditioners, except for triangular factorized forms described in [23]. The drawback of using M that is possibly singular is the need to check the so- lution, or the actual residual norm at the end of the linear iterations. In practice, we have not noticed premature terminations due to a singular preconditioned system, and this is likely a very rare event. We begin this section with an easy proposition. Proposition 3.1. Assume that A is nonsingular and that the residual of the approximate inverse M satisfies the relation consistent matrix norm. Then M is nonsingular. Proof. The result follows immediately from the equality (3. and the well-known fact that if kNk ! 1, then I \Gamma N is nonsingular. We note that the result is true in particular for the Frobenius norm, which, although not an induced matrix norm, is consistent. It may sometimes be the case that AM is poorly balanced and as a result I \Gamma AM can be large. Then balancing AM can yield a smaller norm and possibly a less restrictive condition for the nonsingularity of M . It is easy to extend the previous result as follows. Corollary 3.2. Assume that A is nonsingular and that there exist two nonsingular diagonal matrices D 1 ; D 2 such that consistent matrix norm. Then M is nonsingular. Proof. Applying the previous result to A implies that will be nonsingular from which the result follows. Of particular interest is the 1-norm. Each column is obtained independently by requiring a condition on the residual norm of the form We typically use the 2-norm since we measure the magnitude of the residual I \Gamma AM using the Frobenius norm. However, using the 1-norm for a stopping criterion allows us to prove a number of simple results. We will assume in the following that we require a condition of the form for each column. Then we can prove the following result. Proposition 3.3. Assume that the condition (3.5) is imposed on each computed column of the approximate inverse and let 1. Any eigenvalue of the preconditioned matrix AM is located in the disc 2. If ! 1, then M is nonsingular. 3. If any k columns of M , with k n, are linearly dependent then at least one residual associated with one of these columns has a 1-norm 1. Proof. To prove the first property we invoke Gershgorin's theorem on the matrix each column of R is the residual vector r . The column version of Gershgorin's theorem, see e.g., [30, 17], asserts that all the eigenvalues of the matrix I \Gamma R are located in the union of the disks centered at the diagonal elements and with radius In other words, each eigenvalue must satisfy at least one inequality of the form from which we get Therefore, each eigenvalue is located in the disk of center 1, and radius . The second property is a restatement of the previous proposition and follows also from the first property. To prove the last point we assume without loss of generality that the first k columns are linearly dependent. Then there are k scalars ff i , not all zero such that We can assume also without loss of generality that the 1-norm of the vector of ff's is equal to one (this can be achieved by rescaling the ff's). Multiplying through (3.7) by A yields which gives Taking the 1-norms of each side, we get Thus at least one of the 1-norms of the residuals r must be 1. We may ask the question as to whether similar results can be shown with other norms. Since the other norms are equivalent we can clearly adapt the above results in an easy way. For example, However, the resulting statements would be too weak to be of any practical value. We can exploit the fact that since we are computing a sparse approximation, the number p of nonzero elements in each column is small, and thus we replace the scalar n in the above inequalities by p [18]. We should point out that the result does not tell us anything about the degree of sparsity of the resulting approximate inverse M . It may well be the case that in order to guarantee nonsingularity, we must have an M that is dense, or nearly dense. In fact, in the particular case where the norm in the proposition is the 1-norm, it has been proved by Cosgrove, D'iaz and Griewank [10] that the approximate inverse may be structurally dense, in that it is always possible to find a sparse matrix A for which M will be dense if kI \Gamma AMk 1 ! 1. 14 E. CHOW AND Y. SAAD Next we examine the sparsity of M and prove a simple result for the case where an assumption of the form (3.5) is made. Proposition 3.4. Let assume that a given element b ij of B satisfies the inequality then the element m ij is nonzero. Proof. From the equality Thus, and Thus, if the condition (3.9) is satisfied, we must have us that if R is small enough, then the nonzero elements of M are located in positions corresponding to the larger elements in the inverse of A. The following negative result is an immediate corollary. Corollary 3.5. Let de defined as in Proposition 3.3. If the nonzero elements of are -equimodular in that then the nonzero sparsity pattern of M includes the nonzero sparsity pattern of A \Gamma1 . In particular, if A \Gamma1 is dense and its elements are -equimodular, then M is also dense. The smaller the value of , the more likely the condition of the corollary will be satisfied. Another way of stating the corollary is that we will be able to compute accurate and sparse approximate inverses only if the elements of the actual inverse have variations in size. Unfortunately, this is difficult to verify in advance. 3.2. Case of a nearly singular A. Consider first a singular matrix A, with a singularity of rank one, i.e., the eigenvalue 0 is single. Let z be an eigenvector associated with this eigenvalue. Then, each subsystem (2.3) that is being solved by MR or GMRES will provide an approximation to the system, except that it cannot resolve the component of the initial residual associated with the eigenvector z. In other words, the iteration may stagnate after a few steps. Let us denote by P the spectral projector associated with the zero eigenvalue, by m 0 the initial guess to the system (2.3), and by r the initial residual. For each column j, we would have at the end of the iteration an approximate solution of the form whose residual is The term P r 0 cannot be reduced by any further iterations. Only the norm of can be reduced by selecting a more accurate ffi. The MR algorithm can also break down when Ar j vanishes, causing a division by zero in the computation of the scalar ff j in step 6 of Algorithm 2.3, although this is not a problem with GMRES. An interesting observation is that in case A is singular, M is not too well defined. Adding a rank-one matrix zv T to M will indeed yield the same residual Assume now that A is nearly singular, in that there is one eigenvalue ffl close to zero with an associated eigenvector z. Note that for any vector v we have If z and v are of norm one, then the residual is perturbed by a magnitude of ffl. Viewed from another angle, we can say that for a perturbation of order ffl in the residual, the approximate inverse can be perturbed by a matrix of norm close to one. 3.3. Eigenvalue clustering around zero. We observed in many of our experiments that often the matrix M obtained in a self-preconditioned iteration would admit a cluster of eigenvalues around the origin. More precisely, it seems that if at some point an eigenvalue of AM moves very close to zero, then this singularity tends to persist in the later stages in that the zero eigenvalue will move away from zero only very slowly. These eigenvalues seem to slow-down or even prevent convergence. In this section, we attempt to analyze this phenomenon. We examine the case where at a given intermediate iteration the matrix M becomes exactly singular. We start by assuming that a global MR iteration is taken, and that the preconditioned matrix AM is singular, i.e., there exists a nonzero vector z such that In our algorithms, the initial guess for the next (outer) iteration is the current M , so the initial residual is . The matrix M 0 resulting from the next self- preconditioned iteration, either by a global MR or GMRES step, will have a residual of the form in which is the residual polynomial. Multiplying (3.10) to the right by the eigenvector z yields ae(AM)z As a result we have showing that z is an eigenvector of AM 0 associated with the eigenvalue zero. This result can be extended to column-oriented iterations. First, we assume that the preconditioning M used in self-preconditioning all n inner iterations in a given outer loop is fixed. In this case, we need to exploit a left eigenvector w of AM associated with the eigenvalue zero. Proceeding as above, let m 0 j be the new j-th column of the approximate inverse. We have is the residual polynomial associated with the MR or GMRES algorithm for the j-th column, and is of the form ae j ts j (t). Multiplying (3.11) to the left by the eigenvector w T yields As a result w T Am 0 which can be rewritten as w T AM This gives establishing the same result on the persistence of a zero eigenvalue as for the global iteration. We finally consider the general column-oriented MR or GMRES iterations, in which the self-preconditioner is updated from one inner iteration to the next. We can still write Let M 0 be the new approximate inverse resulting from updating only column j. The residual associated with M 0 has the same columns as those of the residual associated with M except for the j-th column which is given above. Therefore If w is again a left eigenvector of AM associated with the eigenvalue zero, then multiplying the above equality to the left by w T yields showing once more that the zero eigenvalue will persist. 3.4. Convergence behavior of self-preconditioned MR. Next we wish to consider the convergence behavior of the algorithms for constructing an approximate inverse. We are particularly interested in the situation where self-preconditioning is used, but no numerical dropping is applied. 3.4.1. Global MR iterations. When self-preconditioning is used in the global MR iteration, the matrix which defines the search direction is Z is the current residual. Therefore, the algorithm (without dropping) is as follows. 1. R k := I \Gamma AM k 2. Z k := MR k 3. ff k := hRk ;AZk i 4. M k+1 := At each step the new residual matrix R k+1 satisfies the relation Our first observation is that R k is a polynomial in R 0 . This is because, from the above relation, Thus, by induction, in which p j is a certain polynomial of degree j. Throughout this section we use the notation The following recurrence is easy to infer from (3.12), Note that B k+1 is also a polynomial of degree 2 k in B 0 . In particular, if the initial B 0 (equivalently R 0 ) is symmetric, then all subsequent R k 's and B k 's are also symmetric. This is achieved when the initial M is a multiple of A T , i.e., when We are now ready to prove a number of simple results. Proposition 3.6. If the self-preconditioned MR iteration converges, then it does so quadratically. Proof. Define for any ff, Recall that ff k achieves the minimum of kR(ff)k F over all ff's. In particular, This proves quadratic convergence at the limit. The following proposition is a straightforward generalization to the matrix case of a well-known result [13] concerning the convergence of the vector Minimal Residual iteration. Proposition 3.7. Assume that at a given step k, the matrix B k is positive definite. Then, the following relation holds, (R with cos 6 (R; BR) j in which min (B) is the smallest eigenvalue of 1(B +B T ) and oe max (B) is the largest singular value of B. Proof. Start with By construction, the new residual R k+1 is orthogonal to AZ k , in the sense of the h\Delta; \Deltai inner product, and as a result, the second term in the right-hand side of the above equation vanishes. Noting that AZ F F The result (3.16) follows immediately. To derive (3.17), note that in which r i is the i-th column of R, and similarly For each i we have and The result follows after substituting these relations in the ratio (3.17). Note that because of (3.16) the Frobenius norm of R k+1 is bounded from above for all k, specifically, kR k+1 kF kR 0 kF for all k. A consequence is that the largest singular value of B also bounded from above. Specifically, we have oe Assume now that M so that all matrices B k are symmetric. If in addition, each B k is positive definite with its smallest eigenvalue bounded from below by a positive will converge to the identity matrix. Further, the convergence will be quadratic at the limit. 3.4.2. Column-oriented MR iterations. The convergence result may be extended to the case where each column is updated individually by exactly one step of the MR algorithm. Let M be the current approximate inverse at a given sub- step. The self-preconditioned MR iteration for computing the j-th column of the next approximate inverse is obtained by the following sequence of operations. 1. r j := e 2. 3. 4. Note that ff j can be written as where we define to be the preconditioned matrix at the given substep. We now drop the index j to simplify the notation. The new residual associated with the current column is given by We use the orthogonality of the new residual against AMr to obtain kr new k 2 Replacing ff by its value defined above we get kr new k 2 Thus, at each inner iteration, the residual norm for the j-th column is reduced according to the formula kr new in which 6 (u; v) denotes the acute angle between the vectors u and v. Assuming that each column converges, the preconditioned matrix B will converge to the identity. As a result of this, the angle will tend to 6 therefore the convergence ratio sin 6 (r; Br) will also tend to zero, showing superlinear convergence. We now consider equation (3.21) more carefully in order to analyze more explicitly the convergence behavior. We will denote by R the residual matrix We observe that sin 6 This results in the following statement. Proposition 3.8. Assume that the self-preconditioned MR algorithm is employed with one inner step per iteration and no numerical dropping. Then the 2-norm of each residual of the j-th column is reduced by a factor of at least kI \Gamma AMk 2 , where M is the approximate inverse before the current step, i.e., kr new In addition, the Frobenius norm of the residual matrices R obtained after each outer iteration, satisfies As a result, when the algorithm converges, it does so quadratically. Proof. Inequality (3.22) was proved above. To prove quadratic convergence, we first transform this inequality by using the fact that kXk 2 kXkF to obtain kr new Here the k index corresponds to the outer iteration and the j-index to the column. We note that the Frobenius norm is reduced for each of the inner steps corresponding to the columns, and therefore This yields kr new F kr j k 2which, upon summation over j gives This completes the proof. It is also easy to show a similar result for the following variations: 1. MR with an arbitrary number of inner steps, 2. GMRES(m) for an arbitrary m. These follow from the fact that the algorithms deliver an approximate column which has a smaller residual than what we obtain with one inner step MR. We emphasize that quadratic convergence is guaranteed only at the limit and that the above theorem does not prove convergence. In the presence of numerical dropping, the proposition does not hold. 4. Numerical experiments and observations. Experiments with the algorithms and options described in x2 were performed with matrices from the Harwell-Boeing sparse matrix collection [12], and matrices extracted from example problems in the FIDAP fluid dynamics analysis package [16]. The matrices were scaled so that the 2-norm of each column is unity. In each experiment, we report the number of GMRES(20) steps to reduce the initial residual of the right-preconditioned linear system by 10 \Gamma5 . A zero initial guess was used, and the right-hand-side was constructed so that the solution is a vector of all ones. A dagger (y) in the tables below indicates that there was no convergence in 500 iterations. In some tables we also show the value of the Frobenius norm (1.3). Even though this is the function that we minimize, we see that it is not always a reliable measure of GMRES convergence. All the results are shown as the outer iterations progress. In Algorithm 2.4 (dropping in solution vectors) one inner iteration was used unless otherwise indicated; in algorithm 2.5 (dropping in residual vectors) one additional fill-in was allowed per iteration. Various codes in FORTRAN 77, C++, and Matlab were used, and run in 64-bit precision on Sun workstations and a Cray C90 supercomputer. We begin with a comparison of Newton, 'global' and column-oriented iterations. Our early numerical experiments showed that in practice, Newton iteration converges very slowly initially and is more adversely affected by numerical dropping. Global iterations were also worse than column-oriented iterations, perhaps because a single ff defined by (2.2) is used, as opposed to one for each column in the column-oriented case. Table 4.1 gives some numerical results for the WEST0067 matrix from the Harwell-Boeing collection; the number of GMRES iterations is given as the number of outer iterations increases. The MR iteration used self-preconditioning with a scaled transpose initial guess. Dropping based on numerical values in the intermediate solutions was performed on a column-by-column basis, although in the Newton and global iterations this restriction is not necessary. In the presence of dropping we did not find much larger matrices where Newton iteration gave convergent GMRES iterations. Scaling each iterate M i by 1=kAM i k 1 did not alleviate the effects of dropping. The superior behavior of global iterations in the presence of dropping in Table 4.1 was not typical. Table WEST0067: Newton, global, and column MR iterations. dropping Newton y 414 158 100 41 Global 228 102 25 MR Newton 463 y 435 y 457 Global MR 281 120 86 61 43 The eigenvalues of the preconditioned WEST0067 matrix are plotted in Fig. 4.1, both with and without dropping, using column-oriented MR iterations. As the iterations proceed, the eigenvalues of the preconditioned system become closer to 1. Numerical dropping has the effect of spreading out the eigenvalues. When dropping is severe and spoiling occurs, we have observed two phenomena: either dropping causes 22 E. CHOW AND Y. SAAD some eigenvalues to become negative, or some eigenvalues stay clustered around the origin. -0.3 -0.2 -0.10.10.3 (a) no dropping, -0.3 -0.2 -0.10.10.3 (b) no dropping, -0.3 -0.2 -0.10.10.3 (c) -0.3 -0.2 -0.10.10.3 Fig. 4.1. Eigenvalues of preconditioned system, WEST0067 Next we show some results on matrices that arise from solving the fully-coupled Navier-Stokes equations. The matrices were extracted from the FIDAP package at the final nonlinear iteration of each problem in their Examples collection. The matrices are from 2-dimensional finite element discretizations using 9-node quadrilateral elements for velocity and temperature, and linear discontinuous elements for pressure. Table 4.2 lists some statistics about all the positive definite matrices from the collection. The combination of ill-conditioning and indefiniteness of the other matrices was too difficult for our methods, and their results are not shown here. All the matrices are also symmetric, except for Example 7. None of the matrices could be solved with ILU(0) or ILUT [32], a threshold incomplete LU factorization, Table FIDAP Example matrices. Example n nnz Flow past a circular cylinder 7 1633 54543 Natural convection in a square cavity 9 3363 99471 Jet impingement in a narrow channel flow over multiple steps in a channel 13 2568 75628 Axisymmetric flow through a poppet valve of a liquid in an annulus radiation heat transfer in a cavity even with large amounts of fill-in. Our experience with these matrices is that they produce unstable L and U factors in (1.2). Table 4.3 shows the results of preconditioning with the approximate inverse, using dropping in the residual search direction. Since the problems are very ill-conditioned but positive definite, a scaled identity initial guess with no self-preconditioning was used. The columns show the results as the iterations and fill-in progress. Convergent GMRES iterations could be achieved even with lfil as small as 10, showing that an approximate inverse preconditioner much sparser than the original matrix is possible. Table Number of GMRES iterations vs. lfil. 9 203 117 67 51 28 26 24 24 For comparison, we solve the same problems using perturbed ILU factorizations. Perturbations are added to the inverse of diagonal elements to avoid small pivots, and thus control the size of the elements in the L and U factors. We use a two-level block ILU strategy called BILU(0)-SVD(ff), that uses a modified singular value decomposition to invert the blocks. When a block needs to be inverted, it is replaced by the perturbed inverse \Sigma is \Sigma with its singular values thresholded by ffoe 1 , a factor of the largest singular value. Table 4.4 shows the results, using a block size of 4. The method is very successful for this set of problems, showing results comparable to approximate inverse precon- ditioning, but with less work to compute the preconditioner. None of the problems converged, however, for 0:1, and there was not one ff that gave the best result for all problems. We now show our main results in Table 4.5 for several standard matrices in the Harwell-Boeing collection. All the problems are nonsymmetric and indefinite, except for SHERMAN1 which is symmetric, negative definite. In addition, SAYLR3 is singular. SHERMAN2 was reordered with reverse Cuthill-McKee to attempt to change the sparsity pattern of the inverse. Again, we show the number of GMRES iterations to convergence against the number of outer iterations used to compute the approximate inverse. A scaled transpose initial guess was used. When columns in the initial guess contained more than lfil nonzeros, dropping was applied to the guess. Table preconditioner. Example ff= 0.3 ff= 1.0 9 28 72 Numerical dropping was applied to the intermediate vectors in the solution, retaining lfil nonzeros and using no drop tolerance. Table Number of iterations vs. no . Matrix self-preconditioned or unself-preconditioned For problems SHERMAN2, WEST0989, GRE1107 and NNC666, the results become worse as the outer iterations progress. This spoiling effect is due to the fact that the descent property is not maintained when dropping is applied to the intermediate solutions. This is not the case when dropping is applied to the search direction, as seen in Table 4.3. Except for SAYLR3, the problems that could not be solved with ILU(0) also could not be solved with BILU(0)-SVD(ff), nor with ILUTP, a variant of ILUT more suited to indefinite problems since it uses partial pivoting to avoid small pivots [29]. ILUTP also substitutes (10 \Gamma4 times the norm of the row when it is forced to take a zero pivot, where ffi is the drop tolerance. ILU factorization strategies simply do not apply in these cases. We have shown the best results after a few trials with different parameters. The method is sensitive to the widely differing characteristics of general matrices, and apart from the comments we have already made for selecting an initial guess and whether or not to use self-preconditioning, there is no general set of parameters that works best for constructing the approximate inverse. The following two tables illustrate some different behaviors that can be seen for three very different matrices. LAPL0324 is a standard symmetric positive definite 2-D Laplacian matrix of order 324. WEST0067 and PORES3 are both indefinite; WEST0067 has very little structure, while PORES3 has a symmetric pattern. Table 4.6 shows the number of GMRES(20) iterations and Table 4.7 shows the Frobenius norm of the residual matrix against the number of outer iterations that were used to compute the approximate inverse. Table Number of iterations vs. no . Matrix lfil init WEST0067 none A T p 130 none A T u 484 481 y 472 y none I p y y y y y 43 none A T u none I p PORES3 none A T p y y y y y none A T u y y 274 174 116 Table vs. no . Matrix lfil init WEST0067 none A T p 4.43 3.21 2.40 1.87 0.95 none A T u 6.07 6.07 6.07 6.07 6.07 none I p 8.17 8.17 8.17 8.17 8.17 LAPL0324 none A T p 7.91 5.69 4.25 3.12 2.23 none A T u 6.62 4.93 4.00 3.41 3.00 none I p 5.34 4.21 3.53 3.08 2.75 PORES3 none A T p 10.78 9.30 8.25 7.66 7.16 none A T u 12.95 12.02 11.48 10.82 10.20 5. Practical variations and applications. Approximate inverses can be expensive to compute for very large and difficult problems. However, their best potential is in combinations with other techniques. In essence, we would like to apply these techniques to problems that are either small, or for which we start close to a good solution in a certain sense. We saw in Table 4.5 that approximate inverses work well with small matrices, most likely because of their local nature. In the next section, we show how smaller approximate inverses may be used effectively in incomplete block tridiagonal factorizations 5.1. Incomplete block tridiagonal factorizations. Incomplete factorization of block tridiagonal matrices has been studied extensively in the past decade [1, 2, 3, 4, 9, 21, 22], but there have been very few numerical results reported for general sparse systems. Banded or polynomial approximations to the pivot blocks have been primarily used in the past, for systems arising from finite difference discretizations of partial differential equations. There are currently very few options for incomplete 26 E. CHOW AND Y. SAAD factorizations of block matrices that require approximate inversion of general large, sparse blocks. The inverse-free form of block tridiagonal factorization is where LA is the strictly lower block tridiagonal part of the coefficient matrix A, UA is the corresponding upper part, and D is a block diagonal matrix whose blocks D i are defined by the recurrence starting with D The factorization is made incomplete by using approximate inverses rather than the exact inverse in (5.2). This inverse-free form only requires matrix-vector multiplications in the preconditioning operation. We illustrate the use of approximate inverses in these factorizations with Example 19 from FIDAP, the largest nonsymmetric matrix in the collection 259879). The problem is an axisymmetric 2D developing pipe flow, using the two-equation k-ffl model for turbulence. A constant block size of 161 was used, the smallest block size that would yield a block tridiagonal system (the last block has size 91). Since the matrix arises from a finite element problem, a more careful selection of the partitioning may yield better results. In the worse case, a pivot block may be singular; this would cause difficulties for several approximate inverse techniques such as [23] if the sparsity pattern is not augmented. In our case, a minimal residual solution in the null space would be returned. Since the matrix contains different equations and variables, the rows of the system were scaled by their 2-norms, and then their columns were scaled similarly. A Krylov subspace size for GMRES of 50 was used. Table 5.1 first illustrates the solution with BILU(0)-SVD(ff) with a block size of 5 for comparison. The infinity-norm condition of the inverse of the block LU factors is estimated with k(LU) \Gamma1 ek1 , where e is the vector of all ones. This condition estimate decreases dramatically as the perturbation is increased. Table Example 19, BILU(0)-SVD(ff). condition GMRES ff estimate steps 0.500 129. 87 1.000 96. 337 Table 5.2 shows the condition estimate, number of GMRES steps to convergence, timings for setting up the preconditioner and the iterations, and the number of nonzeros in the preconditioner. The method BTIF denotes the inverse-free factorization (5.1), and may be used with several approximate inverse techniques. MR-s(lfil) and MR-r(lfil) denote the minimal residual algorithm using dropping in the solution and residual vectors, respectively, and LS is the least squares solution using the sparsity pattern of the pivot block as the sparsity pattern of the approximate inverse. The MR methods used lfil of 10, and specifically, 3 outer and 1 inner iteration for MR-s, and lfil iterations for MR-r. Self-preconditioning and transpose initial guesses were used. LS used the DGELS routine in LAPACK to compute the least squares solu- tion. The experiments were carried out on one processor of a Sun Sparcstation 10. The code for constructing the incomplete block factorization is somewhat inefficient in two ways: it transposes the data structure of the pivot block and the inverse (to use column-oriented algorithms), and it counts the number of nonzeros in the sparse matrix-matrix multiplication before performing the actual multiplication. Table Example 19, block tridiagonal incomplete factorization. cond. GMRES CPU time est. steps precon solve total precon The timings show that BTIF-MR-s(10) is comparable to BILU(0)-SVD(0.5) but uses much less memory. Although the actual number of nonzeros in the matrix is 259 879, there were 39 355 block nonzeros required in BILU(0), and therefore almost a million entries that needed to be stored. BILU(0) required more time in the iterations because the preconditioner was denser, and needed to operate with much smaller blocks. The MR methods produced approximate inverses that were sparser than the original pivot blocks. The LS method produces approximate inverses with the same number of nonzeros as the pivot blocks, and thus required greater storage and computation time. The solution was poor, however, possibly because the second, third, and fourth pivot blocks were poorly approximated. In these cases, at least one local least squares problem had linearly independent columns. No pivot blocks were singular. 5.2. Improving a preconditioner. In all of our previous algorithms, we sought a matrix M to make AM close to the identity matrix. To be more general, we can seek instead an approximation to some matrix B. Thus, we consider the objective function F in which B is some matrix to be defined. Once we find a matrix M whose objective function (5.3) is small enough, then the preconditioner for the matrix A is defined by This implies that B is a matrix which is easy to invert, or rather, that solving systems with B should be inexpensive. At one extreme when A, the best M is the identity matrix, but solves with B are expensive. At the other extreme, we find our standard situation which corresponds to I , and which is characterized by trivial B-solves but expensive to obtain M matrices. In between these two extremes there are a number of appealing compromises, perhaps the simplest being the block diagonal of A. 28 E. CHOW AND Y. SAAD Another way of viewing the concept of approximately minimizing (5.3) is that of improving a preconditioner. Here B is an existing preconditioner, for example, an LU factorization. If the factorization gives an unsatisfactory convergence rate, it is difficult to improve it by attempting to modify the L and U factors. One solution would be to discard this factorization and attempt to recompute a fresh one, possibly with more fill-in. Clearly, this may be wasteful especially in the case when this process must be iterated a few times due to persistent failures. For a numerical example of improving a preconditioner, we use approximate inverses to improve the block-diagonal preconditioners for the ORSREG1, ORSIRR1 and matrices. The experiments used dropping on numerical values with In Table 5.3, block size is the block size of the block-diagonal preconditioner, and block precon is the number of GMRES iterations required for convergence when the block-diagonal preconditioner is used alone. The number of GMRES iterations is shown against the number of outer iterations used to improve the preconditioner. Table Improving a preconditioner. block block Matrix Besides these applications, we have used approximate inverse techniques for several other purposes. Like in (5.3), we can generalize our problem to minimize F where b is a right-hand side and x is an approximate sparse solution. The right-hand side b does not need to be sparse if dropping is used in the search direction. Sparse approximate solutions to linear systems may be used in forming preconditioners, for example, to form a sparse approximation to a Schur complement or its inverse. See and [8] for more details. 6. Conclusion. This paper has described an approach for constructing approximate inverses via sparse-sparse iterations. The sparse mode iterations are designed to be economical, however, their cost is still not competitive with ILU factorizations. Other approximate inverse techniques that use adaptive sparsity selection schemes also suffer from the same drawback. However, several examples show that these preconditioners may be applied to cases where other existing options, such as perturbed ILU factorizations, fail. More importantly, our conclusion is that the greatest value of sparse approximate inverses may be their use in conjunction with other preconditioners. We demonstrated this with incomplete block factorizations and improving block diagonal pre- conditioners. They have also been used successfully for computing sparse solutions when constructing preconditioners, and one variant has the promise of computing approximations to operators that may be effectively dense. Two limitations of approximate inverses in general are their local nature, and the question of whether or not an inverse can be approximated by a sparse matrix. Their local nature suggests that their use is more effective on small problems, for example the pivot blocks in incomplete factorizations, or else large amounts of fill-in must be allowed. In current work, Tang [33] couples local inverses over a domain in a Schur complement approach. Preliminary results are consistently better than when the approximate inverse is applied directly to the matrix, and its effect has similarities to [7]. In trying to ensure that there is enough variation in the entries of the inverse for a sparse approximation to be effective, we have tried reordering to reduce the profile of a matrix. In a very different technique, Wan et. al. [34] compute the approximate inverse in a wavelet space, where there may be greater variations in the entries of the inverse, and thus permit a better sparse approximation. Acknowledgments . The authors are grateful to the referees for their comments which substantially improved the quality of this paper. The authors also wish to acknowledge the support of the Minnesota Supercomputer Institute which provided the computer facilities and an excellent environment to conduct this research. --R Incomplete block matrix factorization preconditioning methods. On some versions of incomplete block-matrix factorization iterative methods On approximate factorization methods for block matrices suitable for vector and parallel processors Iterative solution of large scale linear systems Iterative solution of large sparse linear systems arising in certain multidimensional approximation problems Approximate inverse techniques for block-partitioned matrices Block preconditioning for the conjugate gradient method Direct Methods for Sparse Matrices Sparse matrix test problems Variational iterative methods for non-symmetric systems of linear equations A stability analysis of incomplete LU factorizations FIDAP: Examples Manual Matrix Computations Parallel preconditioning with sparse approximate inverses Parallel preconditioning and approximate inverses on the Connection Machine The Theory of Matrices in Numerical Analysis On a family of two-level preconditionings of the incomplete block factorization type Modified block-approximate factorization strategies Private communication Efficient parallel solution of linear systems An improved Newton iteration for the generalized inverse of a matrix Preconditioning techniques for indefinite and nonsymmetric linear systems Effective sparse approximate inverse preconditioners. Fast wavelet-based sparse approximate inverse preconditioners --TR --CTR Davod Khojasteh Salkuyeh , Faezeh Toutounian, BILUS: a block version of ILUS factorization, The Korean Journal of Computational & Applied Mathematics, v.15 n.1-2, p.299-312, May 2004 Philippe Guillaume , Yousef Saad , Masha Sosonkina, Rational approximation preconditioners for sparse linear systems, Journal of Computational and Applied Mathematics, v.158 n.2, p.419-442, 15 September Prasanth B. Nair , Arindam Choudhury , Andy J. Keane, Some greedy learning algorithms for sparse regression and classification with mercer kernels, The Journal of Machine Learning Research, 3, 3/1/2003 Kai Wang , Jun Zhang, Multigrid treatment and robustness enhancement for factored sparse approximate inverse preconditioning, Applied Numerical Mathematics, v.43 n.4, p.483-500, December 2002 M. Sosonkina , Y. Saad , X. Cai, Using the parallel algebraic recursive multilevel solver in modern physical applications, Future Generation Computer Systems, v.20 n.3, p.489-500, April 2004 N. Guessous , O. Souhar, Multilevel block ILU preconditioner for sparse nonsymmetric M-matrices, Journal of Computational and Applied Mathematics, v.162 n.1, p.231-246, 1 January 2004 Edmond Chow , Michael A. Heroux, An object-oriented framework for block preconditioning, ACM Transactions on Mathematical Software (TOMS), v.24 n.2, p.159-183, June 1998 T. Tanaka , T. Nodera, Effectiveness of approximate inverse preconditioning by using the MR algorithm on an origin 2400, Proceedings of the third international conference on Engineering computational technology, p.115-116, September 04-06, 2002, Stirling, Scotland Oliver Brker , Marcus J. Grote, Sparse approximate inverse smoothers for geometric and algebraic multigrid, Applied Numerical Mathematics, v.41 n.1, p.61-80, April 2002 Edmond Chow, Parallel Implementation and Practical Use of Sparse Approximate Inverse Preconditioners with a Priori Sparsity Patterns, International Journal of High Performance Computing Applications, v.15 n.1, p.56-74, February 2001 J. Martnez , G. Larrazbal, Wavelet-based SPAI pre-conditioner using local dropping, Mathematics and Computers in Simulation, v.73 n.1, p.200-214, 6 November 2006 Michele Benzi, Preconditioning techniques for large linear systems: a survey, Journal of Computational Physics, v.182 n.2, p.418-477, November 2002

approximate inverse;threshold dropping strategies;preconditioning;krylov subspace methods

292401

Perturbation Analyses for the Cholesky Downdating Problem.

New perturbation analyses are presented for the block Cholesky downdating problem These show how changes in R and X alter the Cholesky factor U. There are two main cases for the perturbation matrix $\D R$ in R: (1) $\D R$ is a general matrix; (2)$\D R$ is an upper triangular matrix. For both cases, first-order perturbation bounds for the downdated Cholesky factor U are given using two approaches --- a detailed "matrix--vector equation" analysis which provides tight bounds and resulting true condition numbers, which unfortunately are costly to compute, and a simpler "matrix equation" analysis which provides results that are weaker but easier to compute or estimate. The analyses more accurately reflect the sensitivity of the problem than previous results. As $X\rightarrow 0$, the asymptotic values of the new condition numbers for case (1) have bounds that are independent of $\kappa_2(R)$ if $R$ was found using the standard pivoting strategy in the Cholesky factorization, and the asymptotic values of the new condition numbers for case (2) are unity. Simple reasoning shows this last result must be true for the sensitivity of the problem, but previous condition numbers did not exhibit this.}

Introduction . Let A 2 R n\Thetan be a symmetric positive definite matrix. Then there exists a unique upper triangular matrix R 2 R n\Thetan with positive diagonal elements such that A = R T R. This factorization is called the Cholesky factorization, and R is called the Cholesky factor of A. In this paper we give perturbation analyses of the following problem: given an upper triangular matrix R 2 R n\Thetan and a matrix X 2 R k\Thetan such that R T is positive definite, find an upper triangular matrix U 2 R n\Thetan with positive diagonal elements such that This problem is called the block Cholesky downdating problem, and the matrix U is referred to as the downdated Cholesky factor. The block Cholesky downdating problem has many important applications, and the case for k=1 has been extensively studied in the literature (see [1, 2, 3, 8, 11, 12, 16, 17, 18]). Perturbation results for the single Cholesky downdating problem were presented by Stewart [18]. Eld'en and Park [10] made an analysis for block downdating. But these two papers just considered the case that only R or X is perturbed. More complete analyses, with both R and X being perturbed, were given by Pan [15] and Sun [20]. Pan [15] gave first order perturbation bounds for single downdating. Sun [20] gave strict, also first order perturbation bounds for single downdating and first order perturbation bounds for block downdating. The main purpose of this paper is to establish new first order perturbation results and present new condition numbers which more closely reflect the true sensitivity of This research was partially supported by NSERC of Canada Grant OGP0009236. y School of Computer Science, McGill University, Montreal, Quebec, Canada, H3A 2A7, ([email protected]), ([email protected]). the problem. In Section 2 we will give the key result of Sun [20], and a new result using the approach of these earlier papers. In Section 3 we present new perturbation results, first by the straightforward matrix equation approach, then by the more detailed and tighter matrix-vector equation approach. The basic ideas behind these two approaches were discussed in Chang, Paige and Stewart [6, 7]. We give numerical results and suggest practical condition estimators in Section 4. Previous papers implied the change \DeltaR in R was upper triangular, and Sun [20] said this, but neither he nor the others made use of this. In fact a backward stable algorithm for computing U given R and X would produce the exact result U for nearby data R+ \DeltaR and X it is not clear that \DeltaR would be upper triangular - the form of the equivalent backward error \DeltaR would depend on the algorithm, and if it were upper triangular, it would require a rounding error analysis to show this. Thus for completeness it seems necessary to consider two separate cases - upper triangular \DeltaR and general \DeltaR. We do this throughout Sections 3-4, and get stronger results for upper triangular \DeltaR than in the general case. In any perturbation analysis it is important to examine how good the results are. In Section 3.2 we produce provably tight bounds, leading to the true condition numbers (for the norms chosen). The numerical example in Section 4 indicates how much better the results of this new analysis can be compared with some earlier ones, but a theoretical understanding is also desirable. By considering the asymptotic case as 0, the results simplify, and are easily understandable. We show the new results have the correct properties as X ! 0, in contrast to earlier results. Before proceeding, let us introduce some notation. Let n\Thetan , then up(B), sut(B), slt(B) and diag(B) are defined by and 2. Previous perturbation results, and an improvement. The condition number for downdating presented by Pan [15] involves the square of the condition number of R, - 2 proposed new condition numbers which are simple and proportional to - 2 (R). The condition number for the block Cholesky downdating problem proposed in [20] is (\Gamma) is the smallest singular value of \Gamma. Notice that for fixed 1. Now we use a similar approach to derive a new bound, from which Sun's bound follows. First we derive some relationships among U , R, X and \Gamma. From (1.1) obviously we have From (1.1) it follows that so that taking the 2-norm gives From (1.1) we have which, combined with (2.2), gives From (1.1) we have which, combined with (2.2), gives s By (2.2) we have Finally from (2.4) we see To derive first order perturbation results we consider the perturbed version of where U , U+ \DeltaU and R are upper triangular matrices with positive diagonal elements. when \DeltaR and \DeltaX are sufficiently small, (2.7) has a unique solution \DeltaU . Multiplying out the two sides of (2.7) and ignoring second order terms, we obtain a linear matrix equation for the first order approximation d \DeltaU to \DeltaU : U T d \DeltaU T In fact it is straightforward to show d U(0), the rate of change of U(-) at so d \DeltaU also has a precise meaning. From (2.8), we have d (R Notice since d \DeltaU U \Gamma1 is upper triangular, it follows, with (1.2), that d (R But for any symmetric matrix B, F \Gamma2 (b 2 Thus from (2.9) we have \DeltaU pkU \GammaT (R T \DeltaR which, combined with (2.2) and (2.3), gives \DeltaU k F - resulting in the new perturbation bound for relative changes \DeltaU k F pp which leads to the condition numbers for the Cholesky downdating problem: pp for U with respect to relative changes in R and X, respectively. Notice from (2.1) . So we can define a new overall condition number Rewriting (2.10) as \DeltaU and combining it with (2.5) and (2.6), gives Sun's bound \DeltaU k F We have seen the right hand side of (2.10) is never worse than that of (2.12), so Although fi 2 is a minor improvement on fi 1 , it is still not what we want. We can see this from the asymptotic behavior of these "condition numbers". The Cholesky factorization is unique, so as X ! 0, U ! R, and X T \DeltaX ! 0 in (2.8). Now for any upper triangular perturbation \DeltaR in R, \DeltaU ! \DeltaR, so the true condition number should approach unity. Here (R). The next section shows how we can overcome this inadequacy. 3. New perturbation results. In Section 2 we saw the key to deriving first order perturbation bounds for U in the block Cholesky downdating problem is the equation (2.8). We will now analyze it in two new approaches. The two approaches have been used in the perturbation analyses of the Cholesky factorization, the QR factorization (see Chang, Paige and Stewart [6, 7]), and LU factorization (see Chang [4] and Stewart [19]). The first approach, the refined matrix equation approach, gives a clear improvement on the previous results, while the second, the matrix-vector equation approach, gives a further improvement still, which leads to the true condition numbers for the block Cholesky downdating problem. 3.1. Refined matrix equation analysis. In the last section we used (2.8) to produce the matrix equation (2.9), and derived the bounds directly from this. We now look at this approach more closely. Let D n be the set of all n \Theta n real positive definite diagonal matrices. For any U . Note that for any matrix B we have First with no restriction on \DeltaR we have from (2.9) d U so taking the F-norm gives \DeltaU It is easy to show for any B 2 R n\Thetan (see Lemma 5.1 in [7]) g. Thus from (3.1) we have \DeltaU (kU \GammaT R T \DeltaR - U (using (2:2); (2:3)) which is an elegant result in the changes alone. It leads to the following perturbation bound in terms of relative changes \DeltaU Although here it would be simpler to just define an overall condition number, for later comparisons it is necessary for us to define the following two quantities as condition numbers for U with respect to relative changes in R and X, respectively (here subscript G refers to general \DeltaR, and later the subscript T will refer to upper triangular \DeltaR): c RG (R; X) where c RG (R; X;D) j Then an overall condition number can be defined as c G (R; X;D); where Obviously we have c G (R; Thus with these, we have from (3.3) that \DeltaU k F if we take become (2.10), and It is not difficult to give an example to show fi 2 can be arbitrarily larger than c G (R; X), as can be seen from the following asymptotic behaviour. It is shown in [7, x5.1, (5.14)] that with an appropriate choice of D, has a bound which is a function of n only, if R was found using the standard pivoting strategy in the Cholesky factorization, and in this case, we see the condition number c G (R; X) of the problem here is bounded independently of - 2 (R) as X ! 0, for general \DeltaR. At the end of this section we give an even stronger result when X ! 0 for the case of upper triangular \DeltaR. Note in the case here that fi 2 in (2.11) can be made as large as we like, and thus arbitrarily larger than c G (R; X). In the case where \DeltaR is upper triangular , we can refine the analysis further. From we have d Notice with (1.3) and (1.4) U \GammaT R T \DeltaRU But for any upper triangular matrix T we have so that if we define T up[diag(U \GammaT R T Thus from (3.12), (3.13) and (3.14) we obtain d \DeltaU As before, let U , where . From (3.15) it follows that \DeltaU U \GammaT \DeltaR T Then, applying (3.2) to this, we get the following perturbation bound \DeltaU k F (3. Comparing (3.16) with (3.3) and noticing (2.3), we see the sensitivity of U with respect to changes in X does not change, so c X (R; X) defined in (3.4) can still be regarded as a condition number for U with respect to changes in X. But we now need to define a new condition number for U with respect to upper triangular changes in R, that is (subscript T indicates upper triangular \DeltaR) c RT (R; X) c RT (R; X;D); where Thus an overall condition number can be defined as where c T (R; Obviously we have With these, we have from (3.16) that \DeltaU k F What is the relationship between c T (R; X) and c G (R; n \Theta n upper triangular matrix observe the following two facts: are the eigenvalues of T , so that which gives (Note: In fact we can prove a slightly sharper inequality ksut(T Therefore c RT (R; n) n) using (2:2)) n)c RG (R; X;D); so that c RT (R; X) n)c RG (R; X): Thus we have from (3.8) and (3.18) n)c G (R; X): On the other hand, c T (R; X) can be arbitrarily smaller than c G (R; X). This can be seen from the asymptotic behaviour, which is important in its own right. As since so for upper triangular changes in R, whether pivoting was used in finding R or not, Thus when X ! 0, the bound in (3.19) reflects the true sensitivity of the problem. For the case of general \DeltaR, if we do not use pivoting it is straightforward to make c G (R; X) in (3.7) arbitrarily large even with 3.2. Matrix-vector equation analysis. In the last subsection, based on the structure of \DeltaR, we gave two perturbation bounds using the so called refined matrix equation approach. Also based on the structure of \DeltaR, we can now obtain provably sharp, but less intuitive results by viewing the matrix equation (2.8) as a large matrix-vector equation. For any matrix C n\Thetan , denote by c (i) j the vector of the first i elements of c j . With this, we define ("u" denotes "upper") c (1)c (2): c (n) It is the vector formed by stacking the columns of the upper triangular part of C into one long vector. First assume \DeltaR is a general real n \Theta n matrix. It is easy to show (2.8) can be rewritten into the following matrix-vector form (cf [7]) WU uvec( d 2 \Theta n(n+1) r 11 and YX 2 R n(n+1) Since U is nonsingular, WU is also, and from (3.22) uvec( d U YX vec(\DeltaX); so taking the 2-norm gives \DeltaU resulting in the following perturbation bound \DeltaU k F -RG (R; X) where Now we would like to show Before showing this, we will prove a more general result. Suppose from (2.8) we are able to obtain a perturbation bound of the form \DeltaU - ff R (3. where ff R and ff X , two functions of R and X, are other measures of the sensitivity of the Cholesky downdating problem with respect to changes in R and X. Let Then from (3.23) and (3.28) we have U ZR vec(\DeltaR)k 2 - ff R Notice \DeltaR can be any (sufficiently small) n \Theta n real matrix, so we must have which gives Similarly, we can show Notice since (3.9) is a particular case of (3.28), (3.27) follows. Thus we have from (3.8) and (3.26) The above analysis shows for general \DeltaR, -RG (R; X) and -X (R; X) are optimal measures of the sensitivity of U with respect to changes in R and X, respectively, and thus the bound (3.24) is optimal. So we propose -RG (R; X) and -X (R; X) as the true condition numbers for U with respect to general changes in R and X, respectively, and - CDG (R; X) as the true overall condition number of the problem in this case. It is easy to observe that if X ! 0, - CDG (R; X) just WU with each entry u ij replaced by r ij . If R was found using the standard pivoting strategy in the Cholesky factorization, then kW \Gamma1 R ZR k 2 has a bound which is a function of n alone (see [5] for a proof). So in this case our condition number - CDG (R; X) also has a bound which is a function of n alone as Remark 1: Our numerical experiments suggest c G (R; X) is usually a good approximation to - CDG (R; X). But the following example shows c G (R; X) can sometimes be arbitrarily larger than - CDG (R; X). where ffl is a small positive number. It is not difficult to show But c G (R; X) has an advantage over - CDG (R; X) - it can be quite easy to estimate - all we need do is choose a suitable D in c G (R; X;D). We consider how to do this in the next section. In contrast - CDG (R; X) is, as far as we can see, unreasonably expensive to compute or estimate. Now we consider the case where \DeltaR is upper triangular. (2.8) can now be rewritten as the following matrix-vector form WU uvec( d 2 \Theta n(n+1) 2 and YX 2 R n(n+1) 2 \Thetakn are defined as before, and WR 2 R n(n+1) 2 \Theta n(n+1) 2 is just WU with each entry u ij replaced by r ij . Since U is nonsingular, WU is also, and from (3.30) uvec( d U YX vec(\DeltaX); so taking the 2-norm gives \DeltaU which leads to the following perturbation bound \DeltaU k F where Note -X (R; X) is the same as that defined in (3.25). As before, we can show that for the case where \DeltaR is upper triangular, - RT (R; X) and -X (R; X) are optimal measures of the sensitivity of U with respect to changes in R and X, respectively, and thus the bound (3.32) is optimal. In particular, we have In fact - RT (R; X) -RG (R; X) can also be proved directly by the fact that the columns of WR form a proper subset of the columns of ZR , and the second inequality has been proved before. Thus we have from (3.8), (3.18), (3.26) and (3.33) By the above analysis, we propose - RT (R; X) and -X (R; X) as the true condition numbers for U with respect to changes in R and X, respectively, and - CDT (R; X) as the true overall condition number, in the case that \DeltaR is upper triangular. If as well X ! 0, then since U ! R, W \Gamma1 So in this case the Cholesky downdating problem becomes very well conditioned no matter how ill conditioned R or U is. Remark 2: Numerical experiments also suggest c T (R; X) is usually a good approximation to - CDT (R; X). But sometimes c T (R; X) can be arbitrarily larger than - CDT (R; X). This can also be seen from the example in Remark 1. In fact, it is not difficult to obtain Like - CDG (R; X), - CDT (R; X) is difficult to compute or estimate. But c T (R; X) is easy to estimate, which is discussed in the next section. 4. Numerical tests and condition estimators. In Section 3 we presented new first order perturbation bounds for the the downdated Cholesky factor U using first the refined matrix equation approach, and then the matrix vector equation approach. We defined - CDG (R; X) for general \DeltaR, and - CDT (R; X) for upper triangular \DeltaR, as the true overall condition numbers of the problem. Also we gave two corresponding practical but weaker condition numbers c G (R; X) and c T (R; X) for the two \DeltaR cases. We would like to choose D such that c G (R; X;D) and c T (R; X;D) are good approximations to c G (R; X) and c T (R; X), respectively. We see from (3.5), (3.6) and (3.17) that we want to find D such that its infimum. By a well known result of van der Sluis [21], - 2 (D \Gamma1 U) will be nearly minimal when the rows of D \Gamma1 U are equilibrated. But this could lead to a large i D . So a reasonable compromise is to choose D to equilibrate U as far as possible while keeping i D - 1. Specifically, take we use a standard condition estimator to estimate Notice from (2.4) we have oe n 2 . Usually k, the number of rows of X, is much smaller than n, so oe n (\Gamma) can be computed in O(n 2 ). If k is not much smaller than n, then we use a standard norm estimator to estimate kXR in O(n 2 ). Similarly kUk 2 and kRk 2 can be estimated in O(n 2 ). So finally c G (R; X;D) can be estimated in O(n 2 ). Estimating c T (R; X;D) is not as easy as estimating c G (R; X;D). The part kdiag(RU (R; X;D) can easily be computed in O(n), since diag(RU c RT (R; X;D) can roughly be estimated in O(n 2 ), based onp F and the fact that kRU \Gamma1 k F can be estimated by a standard norm estimator in O(n 2 ). The value of kXU (R; X;D) can be calculated (if k !! n) or estimated by a standard estimator in O(n 2 ). All the remaining values kRk 2 , kXk 2 and kUk 2 can also be estimated by a standard norm estimator in O(n 2 ). Hence c RT (R; X;D), c X (R; X;D), and thus c T (R; X;D) can be estimated in O(n 2 ). For standard condition estimators and norm estimators, see Chapter 14 of [14]. The relationships among the various overall condition numbers for the Cholesky downdating problem presented in Section 2 and Section 3 are as follows. n)c G (R; X) Now we give one numerical example to illustrate these. The example, quoted from Sun [20], is as follows. 0:240 2:390C C C C C A . The results obtained using MATLAB are shown in Table 4.1 for various values of - : Table c G (R; X;D) 3.60e+03 3.61e+02 3.79e+01 1.79e+01 1.78e+01 1.78e+01 c T (R; X;D) 2.12e+03 2.12e+02 1.79e+01 1.07e+00 1.00e+00 1.00e+00 Note in Table 4.1 how fi 1 and fi 2 can be far worse than the true condition numbers - CDG (R; X) and - CDT (R; X), although fi 2 is not as bad as fi 1 . Also we observe that c G (R; X;D) and c T (R; X;D) are very good approximations to - CDG (R; X) and - CDT (R; X), respectively. When X become small, all of the condition numbers decrease. The asymptotic behavior of c G (R; X;D), c T (R; X;D), - CDG (R; X) and our theoretical results - when - CDG (R; X) will be bounded in terms of n since here R corresponds to the Cholesky factor of a correctly pivoted A, and c T (R; X); - CDT (R; X) ! 1. Acknowledgement . We would like to thank Ji-guang Sun for suggesting to us that the approach describe in (a draft version of) [6] might apply to the Cholesky downdating problem. --R Analysis of a recursive least squares hyperbolic rotation algorithm for signal processing Accurate downdating of least squares solutions A note on downdating the Cholesky factorization PhD Thesis A perturbation analysis for R in the QR factorization New perturbation analyses for the Cholesky factorization Perturbation analyses for the QR factorization Block downdating of least squares solutions Perturbation analysis for block downdating of a Cholesky decomposition Numerical computations for univariate linear models Methods for modifying matrix factor- izations Matrix computations Accuracy and Stability of Numerical Algorithms A perturbation analysis of the problem of downdating a Cholesky factorization Least squares modification with inverse factorizations: parallel implications The effects of rounding error on an algorithm for downdating a Cholesky factor- ization On the Perturbation of LU and Cholesky Factors Perturbation analysis of the Cholesky downdating and QR updating problems Condition numbers and equilibration of matrices --TR

asymptotic condition;downdating;cholesky factorization;perturbation analysis;sensitivity;condition

292410

Condition Numbers of Random Triangular Matrices.

Let Ln be a lower triangular matrix of dimension n each of whose nonzero entries is an independent N(0,1) variable, i.e., a random normal variable of mean 0 and variance 1. It is shown that kn, the 2-norm condition number of Ln, satisfies \begin{equation*} \sqrt[n]{\kn} \rightarrow 2 \:\:\: \text{\it almost surely} \end{equation*} as $n\rightarrow\infty$. This exponential growth of kn with n is in striking contrast to the linear growth of the condition numbers of random dense matrices with n that is already known. This phenomenon is not due to small entries on the diagonal (i.e., small eigenvalues) of Ln. Indeed, it is shown that a lower triangular matrix of dimension $n$ whose diagonal entries are fixed at 1 with the subdiagonal entries taken as independent N(0,1) variables is also exponentially ill conditioned with the 2-norm condition number kn of such a matrix satisfying \begin{equation*} \sqrt[n]{\kn}\rightarrow 1.305683410\ldots \:\:\:\text{\it almost surely} \end{equation*} as $n\rightarrow\infty$. A similar pair of results about complex random triangular matrices is established. The results for real triangular matrices are generalized to triangular matrices with entries from any symmetric, strictly stable distribution.

Introduction Random dense matrices are well-conditioned. If each of the n 2 entries of a matrix of dimension n is an independent N(0; 1) variable, Edelman has shown This work was supported by NSF Grant DMS-9500975CS and DOE Grant DE-FG02- y Department of Computer Science, Cornell University, Ithaca, NY 14853 (diva- [email protected] and [email protected]) Figure 1: Empirical cumulative density functions of n -n , for triangular and unit triangular matrices respectively, with obtained from 1000 random matrices for each n. The random entries are N(0; 1) variables. The higher values of n correspond to the steeper curves. In the limit n ! 1, the cdfs converge to Heaviside step functions with jumps at the dashed lines. that the probability density function (pdf) of -n=n, where -n is the 2-norm condition number of such a matrix, converges pointwise to the function as the distribution of -n=n is independent of n in the limit 1, we can say that the condition numbers of random dense matrices grow only linearly with n. Using this pdf, it can be shown, for example, that E(log(-n In striking contrast, the condition number of a random lower triangular matrix Ln , a matrix of dimension n all of whose diagonal and subdiagonal entries are independent N(0; 1) variables, grows exponentially with n. If -n is the 2-norm condition number of Ln (defined as kLn k show that almost surely as n !1 (Theorem 4.3). Figure 1a illustrates this result. The matrices that arise in the experiments reported in Figure 1 are so ill-conditioned that the standard method of finding the condition number using the SVD [10] fails owing to rounding errors. A numerically stable approach for computing the condition number, which was used to generate the figures, is to find the inverse of the triangular matrix explicitly using the standard algorithm for triangular inversion, and then find the norms of the matrix and its inverse independently. The exponential growth of -n with n is not due to small entries on the diagonal since the probability of a diagonal entry being exponentially small is also exponentially small. For a further demonstration that the diagonal entries do not cause the exponential growth in -n , we consider condition numbers of unit triangular matrices, i.e., triangular matrices with ones on the diagonal. If -n is the condition number of a unit lower triangular matrix of dimension n with subdiagonal entries taken as independent N(0; 1) variables, then almost surely as n !1 (Theorem 5.3). Obviously, in this case the ill-conditioning has nothing to do with the diagonal entries (i.e., the eigenvalues) since they are all equal to 1. Section 7 discusses the relationship of the exponential ill-conditioning of random unit triangular matrices to the stability of Gaussian elimination with partial pivoting. We will use Ln to refer to triangular matrices of various kinds - real or complex, with or without a unit diagonal. But Ln always denotes a lower triangular matrix of dimension n. If the entries of Ln are random variables, they are assumed to be independent. Thus, if we merely say that Ln has entries from a certain distribution, those entries are not only identically distributed but also independent. Of course, only the nonzero entries of Ln are chosen according to that distribution. The condition number always refers to the 2-norm condition number. However, all our results concerning the limits lim n!1 n -n apply to all the L p norms, and the L p norms differ by at most a factor of n. The 2-norm condition number of Ln , defined as denoted by -n . The context will make clear the distribution of the entries of Ln . The analyses and discussions in this paper are phrased for lower, not upper, triangular matrices. However, all the theorems are true for upper triangular matrices as well, as is obvious from the fact that a matrix and its transpose have the same condition number. We obtain similar results for triangular matrices with entries chosen from the complex normal distribution ~ we denote the complex normal distribution of mean 0 and variance oe 2 obtained by taking the real and imaginary parts as independent N(0; oe 2 =2) variables. Let Ln denote a triangular matrix with ~ almost surely as normal entries tend to have smaller condition numbers than triangular matrices with real normally distributed entries. Similarly, let Ln denote a unit lower triangular matrix with ~ subdiagonal entries. Then, almost surely as (Theorem 7.4). Thus, unit triangular matrices with complex normal entries tend to have slightly bigger condition numbers than unit triangular matrices with real normal entries. Our results are similar in spirit to results obtained by Silverstein for random dense matrices [14]. Consider a matrix of dimension n \Theta (yn), where y 2 [0; 1], each of whose n 2 y entries is an independent N(0; 1) variable. Denote its largest and smallest singular values by oe max and oe min , respectively. It is shown in [14] that oe y almost surely as 1. The complex analogues of these results can be found in [4]. The technique used in [14] is a beautiful combination of what is now known as the Golub-Kahan bidiagonalization step in computing the singular value decomposition with the Gerschgorin circle theorem and the Mar-cenko-Pastur semicircle law. The techniques used in this paper are more direct. The exponential growth of is due to the second factor. We outline the approach for determining the rate of exponential growth of -n by assuming Ln triangular with N(0; 1) entries. In Section 2, we derive the joint probability density function for the entries in any column of L \Gamma1 (Proposition 2.1). If T k is the 2-norm of column , i.e., the column with nonzero entries, both positive and negative moments of T k are explicitly derived in Section 3 (Lemma 3.2). These moments allow us to deduce that n converges to 2 almost surely (Theorem 4.3). A similar approach is used to determine the limit of n with ~ (Theorems 5.3, 7.3, and 7.4 respectively). The same approach is used more generally to determine the limit of n as n !1 for Ln with entries drawn from any symmetric, strictly stable distribution (Theorems 8.3 and 8.5). These theorems are specialized to the Cauchy distribution, which is symmetric and strictly stable, in Theorems 8.4 and 8.6. 2 Inverse of a Random Triangular Matrix Consider the matrix ff 11 \Gammaff 21 ff 22 \Gammaff n1 \Gammaff where each ff ij is an independent N(0; 1) variable. Consider L \Gamma1 n , and denote the first k entries in the first column of L \Gamma1 . The t i satisfy the (a) (b) Figure 2: Entries of L \Gamma1 n on the same solid line in (a) have the same pdf. Sets of entries of L \Gamma1 n in the boxes in (b) have the same jpdf. following relations: This system of equations can be interpreted as a system of random recurrence relations. The first entry t 1 is the reciprocal of an N(0; 1) variable. The kth entry t k is obtained by summing the previous entries t independent as coefficients, and dividing that sum by an independent Next, consider an arbitrary column of L \Gamma1 n and denote the first k entries of that column from the diagonal downwards by t . The entries t i satisfy random recurrence relations similar in form to (2.1), but the ff ij are a different block of entries in Ln for different columns. For example, any diagonal entry of n is the reciprocal of an N(0; 1) variable; the kth diagonal entry is 1=ff kk . These observations can be represented pictorially. Every entry of L \Gamma1 n at a fixed distance from the diagonal has the same probability density function (pdf). We may say that the matrix L \Gamma1 n , like Ln , is "statistically Toeplitz." See Figure 2a. Moreover, if we consider the first k entries of a column of L \Gamma1 n from the diagonal downwards, those k entries will have the same joint probability density function (jpdf) irrespective of the column. See Figure 2b. The different columns of L \Gamma1 , however, are by no means independent. Our arguments are stated in terms of the columns of L \Gamma1 n . However, rows and columns are indistinguishable in this problem; we could equally well have framed the analysis in terms of rows. Denote the jpdf of t In the next propo- sition, a recursive formula for f k is derived. For simplicity, we introduce the further notation T k . Throughout this section, Ln is the random triangular matrix of dimension n with N(0; 1) entries. Proposition 2.1. The jpdf f the following recurrence: f Proof. The t k are defined by the random recurrence in (2.1). The expression for f 1 is easy to get. If x is an N(0; 1) variable, its pdf isp The change of variable To obtain the recursive expression (2.3) for f k , consider the variable - k obtained by summing the variables t are independent variables. For fixed values of t the variable being a sum of random normal variables, is itself a random normal variable of mean 0 and variance T 2 . Therefore, the jpdf of - k and t By (2.1), the variable t k can be obtained as - k =ff, where ff is an independent variable. The jpdf of ff, - k and t Changing the variable - k to t integrating out ff, we obtain i.e., f k is given by (2.3). Note that the form of the recurrence for f k in Proposition 2.1 mirrors the random recurrence (2.1) for obtaining t k from the previous entries t In the following corollary, an explicit expression for f k in terms of the t i is stated. Corollary 2.2. For k ? 1, the jpdf f 3 Moments of T k In this section and the next, Ln continues to represent a triangular matrix of dimension n with N(0; 1) entries. As we remarked earlier, the exponential growth of is due to the second factor kL \Gamma1 . Since the 2-norm of column n has the same distribution as Tn\Gammai , we derive formulas for various moments of T k with the intention of understanding the exponential growth of kL \Gamma1 In the lemma below, we consider the expected value E(T - positive and negative values of -. By our notation, j. The notation is used to reduce clutter in the proof. As usual, R k denotes the real Euclidean space of dimension k. The next lemma is stated as a recurrence to reflect the structure of its proof. Lemma 3.2 contains the same information in a simpler form. Lemma 3.1. For any real - ! 1, E(T - k ) is given by the following recurrence: E(T - E(T - E(T - dx For k ) is infinite. Proof. To obtain (3.1), use T and the pdf of t 1 given by Equation (2.2). It is easily seen that the integral is convergent if and only if - ! 1. Next, assume k ? 1. By definition, E(T - Z Using the recursive equation (2.3) for f k , and writing T k in terms of t k and E(T - Z Z R By the substitution t the inner integral with respect to dt k can be reduced to dx Inserting this in the multiple integral (3.3) gives the recursive equation (3.2) for E(T - k ). It is easily seen that the integral in (3.2) is convergent if and only if dx Beginning with the substitution in (3.4), it can be shown that is the beta function. The relevant expression for the beta function B(x; y) is Equation (6.2.1) in [1]. Also, if x is chosen from the standard Cauchy distribution, then We do not need in terms of the beta function, however; the integral expression (3.4) suffices for our purposes. Lemma 3.1 can be restated in a more convenient form using Lemma 3.2. For - for a finite positive constant C - . Also, Proof. The expression for E(T - k ) is a restatement of Lemma 3.1. By elementary and by the form of the integral in (3.4), and Lemma 3.2 implies that the positive moments of T k grow exponentially with k while the negative moments decrease exponentially with k. Obtaining bounds for P now a simple matter. Lemma 3.3. For Proof. \Gamma- to obtain an expression for E(T \Gamma- apply Markov's inequality [2]. Lemma 3.4. For k - 1, Proof. As in Lemma 3.3, - ? 0 implies that P Again, the proof can be completed by obtaining an expression for E(T - using Lemma 3.2 followed by an application of Markov's inequality. 4 Exponential Growth of - n We are now prepared to derive the first main result of the paper, namely, n almost surely as entries. In the sequel, a.s. means almost surely as n ! 1. The definition of almost sure convergence for a sequence of random variables can be found in most textbooks on probability, for example [2]. Roughly, it means that the convergence holds for a set of sequences of measure 1. Lemma 4.1. kLnk 1=n almost surely as n !1. Proof. The proof is easy. We provide only an outline. The Frobenius norm of F , is a sum of n(n variables of mean 1. By the strong law of large numbers, F The proof can be completed using the inequalities n \Gamma1=2 kLnkF - kLn k 2 - kLn The proof of Lemma 4.2 is very similar to the proofs of several standard results in probability, for example the strong law of large numbers [2, p. 80]. Lemma 4.2. As n !1, for any 0 almost surely. Proof. By Lemma 4.1, it suffices to show that a:s: We consider the lower bound first. The 2-norm of the first column of L \Gamma1 n , which has the same distribution as Tn , is less than or equal to kL \Gamma1 . Therefore, for Using Lemma 3.3 with \Gamma- where \Gamma- (fl \Gamma1=- ffl is finite. The first lemma [2] can be applied to obtain infinitely often as n Taking the union of the sets in the above equation over all rational ffl in (0; 1) and considering the complement of that union, we obtain \Gamma- as n In other words, fl \Gamma1=- a.s. The upper bound can be established similarly. At least one of the columns of must have 2-norm greater than or equal to n \Gamma1=2 kL \Gamma1 . Since the 2-norm of column k + 1 has the same distribution as Tn\Gammak , Bounding each term in the summation using Lemma 3.4 gives Lemma 3.2, the largest term in the summand occurs when From this point, the proof can be completed in the same manner as the proof of the lower bound. Theorem 4.3. For random triangular matrices with N(0; 1) entries, as almost surely. Proof. By an inequality sometimes called Lyapunov's [11, p. 144] [2], ff for any real fi ! ff. Thus the bounding intervals [fl \Gamma1=- shrink as - decreases from 1 to 0. A classical theorem [11, p. 139] says that these intervals actually shrink to the following point: lim dx The exact value of the limit can be evaluated to 2 using the substitution tan ' followed by complex integration [3, p. 121]. Thus n Theorem 4.3 holds in exactly the same form if the nonzero entries of Ln are independent N(0; oe 2 ) variables rather than N(0; 1) variables, since the condition number is invariant under scaling. Our approach to Theorem 4.3 began by showing that E(T - - for both positive and negative -. Once these expressions for the moments of T k were obtained, our arguments did not depend in an essential way on how the recurrence was computed. The following note summarizes the asymptotic information about a recurrence that can be obtained from a knowledge of its moments. be a sequence of random variables. If E(jt exponentially with n at the rate - n almost surely as Similarly, if E(jt exponentially with n at the rate - n - as almost surely as n !1. Thus, knowledge of any positive moment of t n yields an upper bound on n knowledge of any negative moment yields a lower bound. 5 Unit Triangular Matrices So far, we have considered triangular matrices whose nonzero entries are inde- pendent, real N(0; 1) variables. In this section and in Section 7, we establish the exponential growth of the condition number for other kinds of random triangular matrices with normally distributed entries. The key steps in the sequence of lemmas leading to the analogs of Theorem 4.3 are stated but not proved. The same techniques used in Sections 2, 3 and 4 work here too. Let Ln be a unit lower triangular matrix of dimension n with N(0; oe 2 ) subdiagonal entries. Let s be the first k entries from the diagonal downwards of any column of L \Gamma1 n . The entries s i satisfy the recurrence are N(0; oe 2 ) variables. The notation S is used below. Proposition 5.1. The jpdf of s by the recur- rence exp(\Gammas 2 exp(\Gammas 2 and the fact that s Lemma 5.2. For any real -, E(S - The note at the end of Section 4 provides part of the link from Lemma 5.2 to the following theorem about -n . Theorem 5.3. For random unit triangular matrices with N(0; oe 2 ) entries, as almost surely. If this limit is denoted by C(oe), then being the Euler constant. Proof. The constant K is given by r- To evaluate K, we used integral 4:333 of [8]. In contrast to the situation in Theorem 4.3, the constant that n to in Theorem 5.3 depends on oe. This is because changing oe scales only the subdiagonal entries of the unit triangular matrix Ln while leaving the diagonal entries fixed at one. For the case discussed in the Introduction, numerical integration shows that the constant is 6 A Comment on the Stability of Gaussian Elimination The conditioning of random unit triangular matrices has a connection with the phenomenon of numerical stability of Gaussian elimination. We pause briefly to explain this connection. For decades, the standard algorithm for solving general systems of linear equations has been Gaussian elimination (with "partial" or row pivot- ing). This algorithm generates an "LU factorization" permutation matrix, L is unit lower triangular with subdiagonal entries - 1 in absolute value, and U is upper triangular. In the mid-1940s it was predicted by Hotelling [12] and von Neumann [9] that rounding errors must accumulate exponentially in elimination algorithms of this kind, causing instability for all but small dimensions. In the 1950s, Wilkinson developed a beautiful theory based on backward error analysis that, while it explained a great deal about Gaussian elimination, confirmed that for certain matrices, exponential instability does indeed occur [17]. He showed that amplification of rounding errors by factors on the order of kL may take place, and that for certain matrices, kL \Gamma1 k is of order 2 n . Thus for certain matrices, rounding errors are amplified by O(2 n ), causing a catastrophic loss of n bits of precision. Despite these facts, the experience of fifty years of computing has established that from a practical point of view, Hotelling and von Neumann were wrong: Gaussian elimination is overwhelmingly stable. In fact, it is not clear that a single matrix problem has ever led to an instability in this algorithm, except for the ones produced by numerical analysts with that end in mind, although Foster and Wright [18] have devised problems leading to instability that plausibly "might have arisen" in applications. The reason appears to be statistical: the matrices A for which kL \Gamma1 k is large occupy an exponentially small proportion of the space of all matrices, so small that such matrices "never" arise in practice. Experimental evidence of this phenomenon is presented in [16]. This raises the question, why are matrices A for which kL \Gamma1 k is large so rare? It is here that the behavior of random unit triangular matrices is rele- vant. A natural hypothesis would be that the matrices L generated by Gaussian elimination are, to a reasonable approximation, random unit triangular matrices with off-diagonal entries of a size dependent on the dimension n. If such matrices could be shown to be almost always well-conditioned, then the stability of Gaussian elimination would be explained. We have just shown, however, that unit triangular matrices are exponentially ill-conditioned. Thus this attempted explanation of the stability of Gaussian elimination fails, and indeed, the same argument suggests that Gaussian elimination should be unstable in practice as well as in the worst case. The resolution of this apparent paradox is that the matrices L produced by Gaussian elimination are far from random. The signs of the entries of these matrices are correlated in special ways that have the effect of keeping kL almost always very small. For example, it is reported in [16] that a certain random matrix A with led to kL was taken to be the same matrix but with the signs of its subdiagonal entries randomized, the result became From a comparison of Theorem 5.3 with half a century of the history of Gaussian elimination, then, one may conclude that unit triangular factors of random dense matrices are very different from random unit triangular matrices. An explanation of this difference is offered in [15] along the following lines. If A is random, then its column spaces are randomly oriented in n-space. This implies that the same holds approximately for the column spaces of L. That condition, in turn, implies that large values kL \Gamma1 k can arise only exponentially rarely. 7 Complex Matrices We now consider matrices with complex entries. Let Ln be a lower triangular matrix with ~ entries. The complex distribution ~ defined in the Introduction . Let t denote the first k entries from the diagonal downwards of any column of L \Gamma1 n . The quantities t k satisfy (2.1), but the ff ij are now independent ~ by R k . Proposition 7.1. The jpdf of r by the recur- rence r 2; (7.1) h for Proof. We sketch only the details that do not arise in the proof of Proposition 2.1. If x and y are independent N(0; oe 2 ) variables, r cos(') and r sin('), then r and ' are independent. Moreover, the distribution of r is Poisson with the pdf Consider the sum - taken as independent ~ variables. For fixed are independent. To see their independence, we write out the equations for Re(- k ) and Im(- k ) as follows: Re(ff ki Re(ff ki )Im(t The linear combinations of Re(ff ki ) and Im(ff ki ) in these two equations can be realized by taking inner products with the two vectors The independence of Re(- k ) and Im(- k ) is a consequence of the orthogonality of v and w, i.e., (v; and the invariance of the jpdf of independent, identically distributed normal variables under orthogonal transformation [13]. Thus for fixed t the real and imaginary parts of - k are independent normal variables of mean 0 and variance R k\Gamma1 =2. By Equation (7.3), the pdfs of are given byR exp(\Gammax=R for positive x; y. The expression (7.2) for h k can now be obtained using r Lemma 7.2. For any - ! 1, E(R - r 2\Gamma- Z 1dx The constant - in Lemma 7.2 can be reduced to ever, as with fl - in Section 3, the integral expression for - is more convenient for our purposes. As before, the note at the end of Section 4 is an essential part of the link from the previous lemma to the following theorem about -n . Theorem 7.3. For random triangular matrices with complex ~ as n !1, almost surely: Theorem 7.3 holds unchanged if the entries are ~ As with Theorem 4.3, this is because the condition number is invariant under scaling. let Ln be a unit lower triangular matrix of dimension n with ~ subdiagonal entries. We state only the final theorem about -n . Theorem 7.4. For random unit triangular matrices with complex ~ tries, as n !1, almost surely, where Ei is the exponential integral. If this limit is denoted by C(oe), then being the Euler constant. Proof. To obtain K, we evaluated using the Laplace transform of log(x) given by integral 4.331.1 of [8]. The explicit formula involving Ei(oe \Gamma2 ) was obtained using integral 4.337.2 of [8]. For -n converges to 8 Matrices with Entries from Stable Distribution The techniques used to deduce Theorem 4.3 require that we first derive the joint density function of the t k , defined by recurrence (2.1), as was done in Proposition 2.1. That proposition made use of the fact that when the ff ki are independent and normally distributed, and the t i are fixed, the sum is also normally distributed. This property of the normal distribution holds for any stable distribution. A distribution is said to be stable, if for X i chosen independently from that distribution, has the same distribution as c n X has the same distribution as are constants [6, p. 170]. If the distribution is said to be strictly stable. As usual, the distribution is symmetric if X has the same distribution as \GammaX . A symmetric, strictly stable distribution has exponent a if c standard result of probability theory says that any stable distribution has an exponent 0 ! a - 2. The normal distribution is stable with exponent The techniques used for triangular matrices with normal entries work more generally when the entries are drawn from a symmetric, strictly stable distri- bution. Let Ln be a unit lower triangular matrix with entries chosen from a symmetric, strictly stable distribution. Denote the pdf of that stable distribution by OE(x). The recurrence for the entries s i of the inverse L \Gamma1 n is given by (5.1), but ff ki are now independent random variables with the density function OE(x). The proposition, the lemma and the theorem below are analogs of Proposition 5.1, Lemma 5.2, and Theorem 5.3 respectively. If the exponent of the stable distribution is a, denote (js 1 j a Proposition 8.1. If OE(x) is the density function of a symmetric, strictly stable distribution with exponent a, the jpdf of s the recurrence and the fact that s Proof. The proof is very similar to the proof of Proposition 2.1. We note that if ff ki are independent random variables with the pdf OE(x), and the s i are fixed, then the sum has the pdf OE(x=S pg. 171]. Lemma 8.2. For any real -, E(S - Theorem 8.3. For random unit triangular matrices with entries from a sym- metric, strictly stable distribution with density function OE(x) and exponent a, as a almost surely: Theorem 5.3 is a special case of Theorem 8.3 when OE(x) is the density function for the symmetric, strictly stable distribution N(0; oe 2 ). Another notable symmetric, strictly stable distribution is the Cauchy distribution with the density function The exponent a for the Cauchy distribution is 1 [6]. Using Theorem 8.3 we obtain, Theorem 8.4. For random unit triangular matrices with entries from the standard Cauchy distribution, as n !1, almost surely. Numerical integration shows the constant to be A similar generalization can be made for triangular matrices without a unit diagonal. However, the analog of Theorem 8.3 for such matrices involves not OE(x), but the density function /(x) of the quotient obtained by taking z as independent variables with the pdf OE. The distribution / can be difficult to compute and work with. let Ln be a triangular matrix with entries chosen from a symmetric, strictly stable distribution with the density function OE(x). We state only the final theorem about -n . Theorem 8.5. For random triangular matrices with entries from a symmetric, strictly stable distribution with density function OE(x) and exponent a, as n !1, a almost surely, where /(x) is the density function of the quotient of two independent variables with the density function OE(x). Theorem 4.3 is a special case of Theorem 8.5 when OE(x) is the density function of the distribution N(0; oe 2 ). The /(x) corresponding to N(0; oe 2 ) is the standard Cauchy distribution. To apply Theorem 8.5 for the Cauchy distribu- tion, we note that log jxj is the density function of the quotient if the numerator and the denominator are independent Cauchy variables. Therefore, Theorem 8.5 implies Theorem 8.6. For random triangular matrices with entries from the standard Cauchy distribution, as n !1, almost surely. The constant of convergence in Theorem 8.6 is 9 Summary Below is a summary of the exponential growth factors lim n!1 n -n that we have established for triangular matrices with normal entries: Real triangular 2 Theorem 4.3 Real unit triangu- lar, Theorem 5.3 Complex triangular e Theorem 7.3 Complex unit trian- gular, Theorem 7.4 The theorems about unit triangular matrices with normally distributed, real or complex entries apply for any variance oe 2 , not just oe of convergence for any symmetric, strictly stable distribution were derived in Theorems 8.3 and 8.5. Those two theorems were specialized to the Cauchy distribution in Theorems 8.4 and 8.6. Similar results seem to hold more generally, i.e., even when the entries of the random triangular matrix are not from a stable distribution. Moreover, the complete knowledge of moments acheived in Lemma 3.2 and its analogs might be enough to prove stronger limit theorems than Theorem 4.3 and its analogs. We will present limit theorems and results about other kinds of random triangular matrices in a later publication. We will also discuss the connection between random recurrences and products of random matrices, and the pseudospectra of infinite random triangular matrices. We close with two figures that illustrate the first main result of this paper, namely, for random triangular matrices with N(0; 1) entries, n almost surely as n !1 (Theorem 4.3). Figure 3 plots the results of a single run of the random recurrence (2.1) to 100; 000 steps, confirming the constant 2 to about two digits. The expense involved in implementing the full recurrence (2.1) for so many steps would be prohibitive. However, since t k grows at the rate 2 k , we need include only a fixed number of terms in (2.1) to compute t k to machine precision. For the figure, we used 200 terms, although half as many would have been sufficient. Careful scaling was necessary to avoid overflow while computing this figure. Figure 4 plots the condition number of a single random triangular matrix for each dimension from 1 to 200. The exponential trend at the rate 2 n is clear, but as in Figure 1, the convergence as n !1 is slow. Acknowledgements We thank D. Coppersmith, P. Diaconis, H. Kesten, A. Odlyzko, J. Sethna, and H. Wilf for helpful discussions. We are especially grateful to Prof. Diaconis for introducing us to stable distributions. Figure 3: Illustration of Theorem 4.3. After 100; 000 steps of the random recurrence has settled to within 1% of its limiting value 2. The implementation is explained in the text. -n Figure 4: Another illustration of Theorem 4.3. Each cross is obtained by computing the condition number -n for one random triangular matrix of dimension n with N(0; 1) entries. The solid line represents 2 n . --R Handbook of Mathematical Func- tions Probability and Measure Functions of One Complex Variable Eigenvalues and Condition Numbers of Random Matrices Eigenvalues and condition numbers of random matrices An Introduction to Probability Theory and Its Applications Gaussian elimination with partial pivoting can fail in practice Table of Integrals Numerical inverting of matrices of high order Matrix Computations Some new methods in matrix calculation Random Matrices and the Statistical Theory of Energy Levels The smallest eigenvalue of a large-dimensional Wishart matrix Analysis of direct methods of matrix inversion A collection of problems for which Gaussian elimination with partial pivoting is unstable --TR --CTR R. Barrio , B. Melendo , S. Serrano, On the numerical evaluation of linear recurrences, Journal of Computational and Applied Mathematics, v.150 n.1, p.71-86, January Kenneth S. Berenhaut , Daniel C. Morton, Second-order bounds for linear recurrences with negative coefficients, Journal of Computational and Applied Mathematics, v.186 n.2, p.504-522, 15 February 2006

exponentially nonnormal matrices;matrix condition numbers;random triangular matrices;strong limit theorems

292837

Continuous Time Matching Constraints for Image Streams.

Corresponding image points of a rigid object in a discrete sequence of images fulfil the so-called multilinear constraint. In this paper the continuous time analogue of this constraint, for a continuous stream of images, is introduced and studied. The constraint links the Taylor series expansion of the motion of the image points with the Taylor series expansion of the relative motion and orientation between the object and the camera. The analysis is done both for calibrated and uncalibrated cameras. Two simplifications are also presented for the uncalibrated camera case. One simplification is made using an affine reduction and the so-called kinetic depths. The second simplification is based upon a projective reduction with respect to the image of a planar configuration. The analysis shows that the constraint involving second-order derivatives are needed to determine camera motion. Experiments with real and simulated data are also presented.

Introduction A central problem in scene analysis is the analysis of 3D-objects from a number of its 2D-images, obtained by projections. In this paper, the case of rigid point con-gurations with known correspondences is treated. The objective is to calculate the shape of the object using the shapes of the images and to calculate the camera matrices, which give the camera movement. One interesting question is to analyse the multilinear constraints that exist between corresponding points in an image sequence. It is well known that corresponding points in two images ful-l a bilinear constraint, known as the epipo- This work has been done within then ESPRIT Reactive LTR project 21914, CUMULI and the Swedish Research Council for Engineering Sciences (TFR), project 95-64-222 lar constraint (Stefanovic 1973, Longuet-Higgins 1981). This constraint is conveniently represented by a 3 \Theta 3 matrix called the essential matrix in the calibrated case and the fundamental matrix in the uncalibrated case. The continuous analogue of this constraint has also been treated in the literature. In this case corresponding representations involve the so called in-nitesimal epipole or the focus of expansion and the axis of rotation. This has been studied by photogrammetrists in the calibrated case and recently by Faugeras and Vieville in the uncalibrated case, cf. (Vieville and Faugeras 1995). The continuous time analogue can be derived from the bilinear constraint as the limit when the time dioeerence tends to zero. Since the bilinear constraint involves two time instants, the continuous time analogue involves the Taylor expansion to the -rst order, the so called one-jet. 2 -str-m and Heyden The goal of this work is to emphasise the connection between the discrete time constraints and corresponding continuous time constraints, and to derive the continuous time analogue of multilinear constraints. The multilinear camera-image motion constraints have been treated in several recent conference pa- pers, (Faugeras and Mourrain 1995, Heyden and #str#m 1995, Triggs 1995). These derivations involve dioeerent types of mathematical techniques and are represented using dioeerent mathematical objects, e.g. Grassman-Cayley algebra, Joint Grassmannian etc. In this paper we will use a matrix formulation, such that the image-motion constraint will be of the type that the rank of a certain matrix is less than full, i.e. that all subde- terminants of this matrix are zero. The elements of this matrix will depend on the motion of the image point and the relative motion and orientation between the camera and the object con-gu- ration as expressed by the so called camera projection matrix. There are other ways of formulating the same constraint. However, this particular choice of formulation has some nice advantages: ffl Using this formulation, it is easier to show that all multilinear constraints can be derived from the bi- and trilinear constraints. ffl The dioeerence and the similarities between the uncalibrated and calibrated case is clearer ffl The parameters in the multilinear constraints are closely linked to the camera parameters describing the projection of points in 3D onto each image plane. ffl The continuous time case can easily be derived from the discrete time case. In this paper we introduce and study the continuous time analogue of these multilinear constraints. The nth order constraint involve the Taylor expansion up to order n of the image point motion and a similar expansion for the motion of the camera relative to the scene. These constraints have the same applications as their discrete counterparts, i.e. they can be used as matching constraints to -nd points that are moving rigidly with respect to the camera. They can also be used to calculate relative motion with no a priori knowledge about the scene. We would like to emphasise the derivation of these constraints. Some experimental validation has been included, but needless to say, more experiments are needed in order to validate the potential of these types of constraints. The paper is organised as follows. Some basic notations are introduced in Section 2. Sections 3, 4 and 5 contain background material on the ambiguity in the choice of coordinate system, camera matrix parametrisation and a short derivation of the discrete time constraints. The continuous time analogue of these constraints is derived in Section 6. These constraints are studied in Sections 7, 8 and 9 with respect to coordinate ambiguity, motion observability and motion esti- mation. The latter is illustrated with simulated and real data in Section 10. Then follows a short discussion and conclusions in Section 11. 2. Camera Geometry and Notation The pinhole camera model is used and formulated using projective geometry. A point in three dimensions with Euclidean co-ordinates represented as a point in the three-dimensional projective space using homogeneous coordinates \Theta U x U y U z 1 T . In projective geometry two representations are considered as the same point if one is a multiple of the other, (Faugeras 1993). Projection onto the image plane is conveniently represented in the camera projection matrix for- mulation where - is the unknown depth and u is the image position, also in homogeneous coordinates \Theta possibly corrected for the internal calibration if this is known. A priori knowledge about the camera and the camera-object motion, give information about the structure of the camera projection matrix P . \Theta I uncalibrated, (2) \Theta I known A, (3) \Theta I known \Theta T denotes the unknown position of the camera focus, R a rotation matrix describing the orientation of the camera, A a matrix representing the unknown internal calibration parameters and I the 3 \Theta 3 identity matrix. Thus in the three situations above there are 11, 6, and 3 degrees of freedom respectively in each camera matrix. In the uncalibrated case when the internal calibration matrix A is unknown it is convenient to Continuous Time Matching Constraints for Image Streams 3 of the camera. In other words we think of the camera as having a position T and a generalised orientation - Q. The position determines how object points are projected onto the viewing sphere and the orientation - rearranges these directions. Both orientation and position of the camera will change over time. In the continuous time case we will use the notation ( - for the orientation and position of the camera at time t, . In the discrete time case the camera at time t will be represented by ( - Z. The image position u(t) of a point U at time t is thus \Theta - Alternatively, the notation and T will be used. The projection is then expressed as I Capital U is used to denote object points. Corresponding lower case letter u is used to denote the corresponding image point. Subscripts, e.g. are used to denote dioeerent points. The superscripts with parenthesis are used to denote co- eOEcients of the Taylor series expansion, i.e. Boldface 0 denotes a zero matrix, usually 3 \Theta 1 or 2 \Theta 1. 3. Uniqueness of Solutions The overall goal is to calculate the camera matrices P (t) and to reconstruct the coordinates of the objects U given only the image positions u(t), such that the camera equation is ful-lled. The camera motion is assumed to be smooth. The solution for P (t) and U can only be given up to an unknown choice of coordinate system in the world. This is easiest seen in the uncalibrated case. Let B be any non-singular 4\Theta4 matrix. Change the projection matrices according to and change object points according to Now U) is another solution with dioeerent coordinates that also ful-l the camera equation For the case of uncalibrated cameras there are 15 degrees of freedom in choosing a projective co-ordinate system. For the case of known internal calibration there is an arbitrary choice of orientation, origin and scale (7 degrees of freedom). In the case of known external and internal orientation we are only allowed to make changes of the type \Deltay i.e. one may choose the origin arbitrarily and also the overall scale. It is important to keep this problem of non-uniqueness in mind. The question of choosing a particular or canonical choice of coordinate sy- stem, as discussed in (Luong and Vieville 1994), can be important when designing numerical algo- rithms. This is commented further upon in Section 7. 4. Simplifications in the Uncalibrated Case In the uncalibrated case we can simplify the problem by partially -xing the object and image co-ordinate systems. Two such simpli-cations are of special interest. These are explained in detail in (Heyden and #str#m 1995, Heyden and #str#m 1997), and will be brieAEy described here. 4.1. Affine Reduction If the same three points can be seen in a sequence of images, they can be used to simplify the problem according to the following theorem. Theorem 1. Let u 1 (t), u 2 (t) u 3 (t) be the images of three points U 1 and U 3 . Choose an object coordinate system where U Choose image coordinate systems at time t so that u 1 Then the camera projection matrices can be written Q(t) is a diagonal matrix and T (t) is the position of the camera at time t. 4 -str-m and Heyden Proof: The choice of coordinate systems gives three constraints on the camera matrix P (t), Since P (t)U j is the jth column of the matrix P (t), it follows that the -rst three columns of P (t) form a diagonal matrix, i.e. By aOEne normalisation of three corresponding points in an image sequence, the analysis of the remaining points can be made almost as if the cameras were calibrated and with the same rotation. The unknown elements in the diagonal matrices correspond to the so called kinetic depths of the three image points relative to the camera motion, cf. (Sparr 1994, Heyden 1995). This is basically the same idea as the relative aOEne coordinates in (Shashua and Navab 1996). Using projective geometry one can think of this as de-ning the plane through the three points as the plane at in-nity and also partially locking the orientation of this plane by these three points. 4.2. Projective Reduction Further simpli-cations can be obtained if four or more coplanar points, e.g. points belonging to a planar curve, are detected in a subsequence of images. Theorem 2. Let u 1 (t), u 2 (t), u 3 (t) and u 4 (t) be the images of four coplanar points U 1 and U 4 so that no three of them are collinear. Choose an object coordinate system where U and U Choose image coordinate systems at time t so that u 1 Then the camera projection matrices can be written \Theta I , where T (t) is the position of the camera at time t. Proof: The -rst 3 \Theta 3 block matrix - Q(t) of P (t) acts as the identity on i.e. where - denotes equality up to scale. According to the assumptions, the four points constitute a projective basis for the projective plane. Therefore, it follows that - I . By projective alignment of the images of at least four coplanar points the problem can thus be analysed as if both internal calibration and orientation of the camera are known at all times. The motion of the camera can be described by the pair (Q(t); T (t)) or ( - (or - describes the generalised orientation of the camera and T (t) (or - describes the position of the camera at time t. Depending on which assumptions and simpli-cations we have made, the matrices Q(t) (or - Q(t)) lie on dioeerent manifolds ffl Traditional uncalibrated setting: The matrices Q(t) are arbitrary but nonsingular and two matrices are considered equal if one is a (positive) multiple of the other. There are eight degrees of freedom. AOEnely reduced setting: The matrices Q(t) are diagonal and nonsingular and two matrices are considered equal if one is a (positive) multiple of the other. There are two degrees of freedom. Projectively reduced setting: The matrices Q(t) are identity matrices. There are no degrees of freedom. Alternatively we may require: The matrices Q(t) are multiples of the identity matrix and two matrices are considered equal if one is a (positive) multiple of the other. Calibrated setting: The matrices Q(t) are ort- hogonal. There are three degrees of freedom. Alternatively we may require: The matrices are multiples of orthogonal matrices and two matrices are considered equal if one is a (positive) multiple of the other. These manifolds are non-linear and have a so called Lie group structure under matrix multipli- cation. The corresponding Lie algebra will be of importance. Unlike the Lie group the Lie algebra is a linear subspace of all 3 \Theta 3 matrices q: ffl Traditional uncalibrated setting: The matrices q(t) are arbitrary and two matrices are considered equal if their dioeerence is a multiple of the identity matrix. There are eight degrees of freedom. AOEnely reduced setting: The matrices q(t) are diagonal and two matrices are conside- Continuous Time Matching Constraints for Image Streams 5 red equal if their dioeerence is a multiple of the identity matrix. There are two degrees of Projectively reduced setting: The matrices q(t) are zero matrices. There are no degrees of freedom. Calibrated setting: The matrices q(t) are an- tisymmetric. There are three degrees of freedom These Lie algebras are obtained from the Lie groups using the exponential map The dioeerent Lie groups and Lie algebras are summarised in Table 1. Since two matrices in the Lie Group are considered to be equal if one is a multiple of the other, it is often convenient to choose a speci-c representative. One such choice of representative is to always scale the matrix so that the determinant is one. Similarly two matrices in the Lie algebra are considered to be equal if the dioeerence is a multiple of the identity. A unique representative can be chosen by demanding that the trace of the matrix is zero. This -ts in nicely with the exponential map since where we have used the notation det as the determinant and tr as the trace of a matrix. 5. Multilinear Constraints in the Discrete Time Case In order to understand the matching constraints in the continuous time case, it is necessary to take a look at the corresponding constraints in the discrete time case. For a more thorough tre- atment, see (Heyden and #str#m 1995). Alternative formulations of the same type of constraints can be found in (Faugeras and Mourrain 1995, Triggs 1995). We start with the de-nition. De-nition 1. The nth order discrete multilinear constraint is Table 1. The Lie-Groups and their corresponding Lie- Algebras in the four dioeerent settings. Setting Lie Group Lie Algebra Uncalibrated Arbitrary Arbitrary Aoe. red. Diagonal Diagonal Proj. red. Calibrated Rotation Antisymmetric Theorem 3. In a sequence of discrete images corresponding points with coordinates u(t 0 obey the nth order discrete multilinear constraint. This means that there exist a solution to (8) for that holds for every corresponding point sequence. Proof: This can be seen by lining up each projection constraint in a linear matrix equation66P (t 0 \GammaU Since this system has a non-trivial solution the leftmost matrix cannot have full rank. 6. Continuous Time Analogue of Multilinear Constraints The multilinear constraints in the continuous time case can be derived using Taylor series expansion of the time dependent functions in (5), Use the Taylor series expansions Recall that superscript (i) denotes the ith coeOE- cient in the Taylor series expansion, e.g. - (k) =k! . Substituting the Taylor series expansions 6 -str-m and Heyden into (10) gives Identifying the coeOEcients of t i for \GammaU Since this set of equations has a non-trivial solution the leftmost matrix cannot have full rank. De-nition 2. The nth order continuous constraint is where Theorem 4. In a continuous sequence of images the image coordinates u (0) and their derivatives, up to order n at the same instant of time obey the nth order continuous constraint. Proof: Follows from the derivation above. Remember that - , and similarly for the other variables. This means in particular that dt has the meaning of image point velocity. The continuous time constraints can be simpli- -ed somewhat by partially choosing a coordinate system according to P (which implies By multiplying the big matrix M from the right by the matrix we obtain MS Since the matrix S has full rank, it follows that By elimination of the -rst three rows and columns of the matrix MS in (16), the constraint (15) is simpli-ed to De-nition 3. The continuous analogue of the bilinear and trilinear constraints are de-ned as the nth order constraint above with respectively. The analogue of the bilinear constraint is the -rst order continuous constraint rank \Theta - and analogue of the trilinear constraint is the second order continuous constraint rank 7. Choice of Coordinate System In the previous sections we have used the option to partially -x the coordinate system to simplify the problem. Sometimes it is useful to choose a canonical object coordinate frame to obtain a canonical coordinate representation of the reconstructed object and projection matrices. In the calibrated case a speci-c coordinate system can be determined by setting the initial orientation to Q(t 0 overall scale by jT (1) (t 0 In the uncalibrated case there are four things to consider when choosing a projective coordinate system in the reconstruction. 1. The position of the plane at in-nity, 3 d 2. The individual points at the plane at in-nity, 3. The origin, 3 d 4. The scale, 1 d Continuous Time Matching Constraints for Image Streams 7 One way of doing this in the discrete time case is to lock individual points at the plane at innity by lock the origin by T and the scale by jT 1. The position of the plane at in-nity can be chosen by choosing a speci-c Q(t 1 ). This cannot, however, be done arbitrarily. The matrix Q(t 1 ) ful-lls the bilinear constraint, as discussed in (Luong and Vieville 1994). The question of choosing a canonical coordinate system (and thereby choosing a speci-c plane at in-nity) is simpler in the aOEnely and projectively reduced settings. The plane at in-nity is determined by the three or more points that are used in the reduction. A canonical coordinate system can then be chosen by Q(t 0 similar to the calibrated case. Similarly in the continuous time case one possible choice is to take P (which implies discussed in the simpli-cation above, and then take j - tr - This determines the choice of coordinate system in the calibrated, aOEnely reduced and in the projectively reduced settings. In the uncalibrated case there is a further choice of the plane at in-nity. This choice can be made by choosing one - that ful-lls the -rst order continuous constraint. Since Q(t) is undetermined up to a scalar factor it is possible to enforce uniqueness if we require det t. This condition is of course ful-lled for I . A Taylor series expansion of Q(t) gives which can be seen by expanding the determinant. Thus we have tr Q The coeOEcients of t k in (20) are complicated expressions in Q (i) . Ho- wever, it can be seen that the coeOEcient of t k can be written as tr Q (k) plus terms involving Q (i) for k. This means that we can ensure uniqueness if we require The price we have to pay for this simpli-cation is that det Q(t) depends on t, det in general for Assume again that Q (0) has been chosen to be the unit matrix. It also follows from the matrix exponential that Q has the properties listed in Table 1. However, the relations between q (i) and Q (i) for i ? 1 are more complicated. 8. Motion Observability What can be said about the observability of the full motion of the camera, e.g. is it possible to determine camera motion uniquely up to a choice of coordinate system using only the -rst order continuous constraint? 8.1. Motion Observability from the First Order Continuous Constraint Does the -rst order continuous constraint determine camera motion uniquely up to choice of camera system? Using the -rst order continuous constraint we can calculate T (1) (t) up to an unknown scale factor and Q (1) (t) up to an arbitrary choice of the plane at in-nity. Since only the direction of T (1) (t) can be obtained at each time instant, it is not possible to reconstruct T uni- quely. Thus, the -rst order continuous constraint is not enough to determine motion. 8.2. Motion Observability from the Second Order Continuous Constraint If T (1) (t) and Q (1) (t) are known and T (1) (t) 6= 0, then T (2) (t) and Q (2) (t) are uniquely determined by the second order continuous constraint (19). One can think of this as T (2) (t) and Q (2) (t) being a function of T (1) (t), Q (1) (t), T (t), Q(t) and image motion at this time instant, i.e. It is a well known fact that this kind of dioee- rential equations can be solved at least locally, given a set of initial conditions on T These initial conditions are determined by choosing a coordinate system and at the same time ful-lling the -rst order continuous constraint at Thus the full motion of the camera is observable from the second order continuous constraints if T (1) (t) 6= 0. Since the 8 -str-m and Heyden full camera motion (T (t); Q(t)) can be calculated uniquely up to a choice of coordinate system, it is also possible to calculate derivatives of all or- ders. Thus all continuous multilinear constraints follow from the -rst and second order continuous constraints. 9. Estimation of Motion Parameters using the Continuous Time Constraints Study again the simpli-cation of the continuous time constraint in (17). One use of this constraint is to calculate ( - the motion of the points in the image u (i) . Typically, the relative noise increases with increasing orders of dioeeren- tiation. We therefore expect u (2) to be noisier than u (1) which in turn is noisier than u (0) . It is therefore natural to estimate motion parameters in steps. First estimate ( - using the -rst order continuous constraint (18). Then try to estimate ( - and the second order continuous constraint (19). Repeat this as long as the level of noise permits. There are some numerical diOEculties with using the continuous time constraint to estimate motion parameters. First of all it can be quite diOEcult to obtain good estimates of the image point positions and their derivatives. Secondly, since these estimates are noisy, this needs to be modelled and taken into account. This aoeects the way motion parameters should be estimated. 9.1. Using the First Order Continuous Constraint The -rst order continuous constraint (18) involves the -rst order derivative in camera motion, T (1) and Q (1) . One major dioeerence between the discrete time case and the continuous time case is that the derivatives of the orientation, Q (1) , lie on a linear manifold. As an example take the calibrated case. Whereas in the discrete case the motion parameters involve a rotation matrix Q(t 1 ), the continuous time analogue involve an anti-symmetric matrix Q (1) . It is easier to parametrise the set of anti-symmetric matrices (since this is a linear subspace of all matrices), than to parametrise the set of rotation matrices (which is a non-linear ma- nifold). The velocity T (1) also lies on a linear manifold. It can, however, only be determined up to scale. If T (1) is known, the problem of determining Q (1) is linear so it can quite easily be solved with linear methods. On the other hand, if Q (1) is known the problem of determining T (1) can be formulated and solved in a linear fashion. This suggests a fast two-step iterative method. Guess Q (1) . Holding -xed solve for T (1) . Holding T (1) -xed solve for Q (1) . This method has been tried and in most cases it does seem to converge nicely. It should be noted, however, that there are no guarantees that this method will converge. An alternative and more robust method is to tessellate the sphere of directions T (1) 2 S 2 . For each T (1) , solve for Q (1) linearly and store the residual. Choose the pair (T (1) ; Q (1) ) that gave the lowest residual. Any of these two methods can be re-ned by taking the motion parameters as an initial estimate of (T (1) ; Q (1) ), in a non-linear least squares minimisation. An advantage of this re-nement is that it allows more sophisticated error measures, e.g. the maximum likelihood estimate, that takes into account the quality of the estimates of u and u (1) . Another advantage of this re-nement is that it allows for an analysis of the stability of the solution through the analysis of the residuals at the optimum. 9.2. Using the Second Order Continuous Constraint The second order continuous constraint in (19) involve the -rst and second order derivatives of the camera motion. As for the -rst order continuous constraint, it should be possible to use a two step iterative method. Guess ( - Holding -xed solve for ( - T (2) ). Holding -xed solve for ( - should be noted that the convergence properties of these methods haven't been studied. Another approach might by to estimate using the -rst order continuous constraint and while holding these -xed, estimate using the second order continuous constraint. Close to a solution the method could be rened by non-linear maximisation of a Likelihood function. 10. Experiments To illustrate the continuous constraints, we have used iterative methods as described above. We consider the -rst order continuous constraint in 3D to 2D projection and the second order constraint in an industrial application of 2D to 1D projection. Continuous Time Matching Constraints for Image Streams 9 10.1. First Order Constraints on Image Streams An image sequence of an indoor scene have been used, see Figure 1, where one image in the sequence is shown. The whole sequence contains more than 200 images. To illustrate the applicability of the continuous constraints, we have only used 2 images. Points have been extracted using a corner detector, and we have used 28 points with correspondences. The aOEnely reduced coordinates have been calculated from the images, giving u(0) and u(h), where h denotes the time increment between the dioeerent exposures. In this case We have used u u(0). The derivatives have been computed from image 1 and image 2 using the dioeerence approximation Using the iterative approach and the aOEnely reduced setting and only 10 iterations, starting from Q we obtain the following solution ful-lling the -rst order constraint: This solution can be compared to the solution obtained in the discrete time case between u(0) and Here - Q(h) should be compared to - I and - T (h) to - T (1) . The angle between - T (1) to - T (h) is 2:8 degrees. 10.2. Second Order Constraints on Angle Streams The continuous multilinear constraints have an interesting practical application to the vision system of an autonomous vehicle. The vehicle is equipped with a calibrated camera with a one-dimensional retina. It can only see speci-c beacons or points in a horizontal plane. Let (x; y) be the coordinates of such a beacon and let as Fig. 1. One image in the sequence used in the continuous case. before this point be represented by a vector \Theta T . Then the measured image point is a direction vector according to the camera equation where the camera matrix P (t) is a 2 \Theta 3 calibrated camera matrix Let the Taylor expansion of the rotation matrix R(t) and the vector - T (t) be ReAEector Angle meter a Fig. 2. a: A laser guided vehicle. b: A laser scanner or angle meter. Using this notation the continuous time analogue of the trilinear constraint (19) becomes det det The matrix R is a 2\Theta2 rotation matrix. Assuming that the Taylor expansion of the rotation angle OE(t) is -rst derivative of R has the form r 1- and the second coeOEcient of the Taylor expansion has the form R The second order continuous constraint in (26) does not, however, involve the term r 2in R (2) . This can be seen by adding r 2times the last column to the second column in (26). The second order continuous constraint in (26) will now be studied in more detail. Introduce the variables a 1 a 2 and The determinant of M can be written \Theta a 1 a 2 22 where \Gammay 0 \Theta a 1 a 2 and \Theta T . A tentative algorithm to -nd w and z has been investigated. Algorithm 1 1. Start with a crude estimate of r 1 and r 2 , for example \Theta T . 2. For all image directions u i , calculate the corresponding . The vector w should be orthogonal to every vector , the vector w can be found as the left null space of the matrix \Theta The vector w is found by using a singular value decomposition of M . 3. Once w is approximately known, z can be found as the right null space of6wV 1 4. Repeat steps 2 and 3. The algorithm has been implemented and tested experimentally. The results are illustrated with simulated data. In these simulations the angles to -fteen beacons were calculated during a period of one second. Gaussian noise of dioeerent standard deviations 0:1; 0:5; added to the angle measurements. In the real application the standard deviation is approximately 0:2 mrad. The angle measurements in a two second period were used to estimate the Taylor co- eOEcients of the measured image direction u i using standard regression techniques. These Taylor co- eOEcients were then used to calculate the Taylor coeOEcients of the motion (r using the algorithm above. The experiment was repeated 100 times. The standard deviation of the estimated motion parameters are shown in Table 2. The true values of the motion parameters, in the simulation are r Continuous Time Matching Constraints for Image Streams 11 Table 2. Standard deviation of estimated motion parameters for dioeerent noise levels when using Algorithm 1. 11. Discussion and Conclusions In this paper the simpli-ed formulation of the multilinear forms that exist between a sequence of images has been used to derive similar constraints in the continuous time case. The new formulation makes it easier to analyse the matching constraints in image sequences. It has been shown that these constraints contain information in the -rst and second order only. This representation is fairly close to the representation of the motion and it is easy to generalise to dioeerent settings. Four such settings calibrated, uncalibrated, aOE- nely reduced and projectively reduced, are described in the paper. The continuous constraints can be used to design -lters to estimate structure and motion from image sequences. Using only the -rst order constraint it is possible to estimate the direction of the movement of the camera. Having only this information it is not possible to build up the whole camera movement. Using the second order constraint it is possible to estimate the second order derivative of the camera movement with a scale consistent with the -rst order deri- vative. This information can be used to build up the camera movement. This is analogous to the discrete time case, where the trilinearities are needed in order to estimate the camera movement, if only multilinear constraints between consecutive images are used. The -rst example above shows that the -rst order continuous constraint is comparable to the discrete case. However, the continuous constraints are sensitive to noise because estimating the image velocities ampli-es the noise present in the images. The higher order continuous constraints are even more sensitive, because they involve higher order derivatives. Using -ltration techniques to estimate the derivatives from image coordinates in more than two images could reduce the in- AEuence of noise. The second example illustrates the applicability of higher order constraints in an industrial application. An industrial vehicle is guided by a one-dimensional visual system. It is interesting to note that the -rst order constraint in this case gives no information about the motion. Higher order constraints are essential. We believe that the continuous constraints are important theoretical tools. We also believe that they are necessary to analyse image sequences with high temporal sampling frequency. Acknowledgements The authors thank Tony Lindeberg and Lars Bretzner at KTH for help with corner detection and tracking. --R Workshop on Intelligent Autonomous Vehicles on Computer Vision the velocity case Topological and Geometrical Aspects of PhD thesis Department of Medical and Physiological Physics What can be seen in three dimensions with an uncalibrated stereo rig? On the geometry and algebra of the point and line correspon between Reconstruction from image sequences by means of relative depths Computer for sequences of images of Visual Scenes forms for sequences of images. in Image and Vision Computing. A computer algorithm for reconstructing a scene from two projections Canonic representations for the geometries of multiple projective views Conf. on Computer Vision Relative aOEne structure: Canonical model for 3d from 2d geometry and applications Machine Intell A common framework for kinetic depth reconstruction and motion for deformable objects Motion analysis with a camera with unknown and possibly varying --TR

multilinear constraint;calibrated camera;structure;motion;optical flow;uncalibrated

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The Hector Distributed Run-Time Environment.

AbstractHarnessing the computational capabilities of a network of workstations promises to off-load work from overloaded supercomputers onto largely idle resources overnight. Several capabilities are needed to do this, including support for an architecture-independent parallel programming environment, task migration, automatic resource allocation, and fault tolerance. The Hector distributed run-time environment is designed to present these capabilities transparently to programmers. MPI programs can be run under this environment on homogeneous clusters with no modifications to their source code needed. The design of Hector, its internal structure, and several benchmarks and tests are presented.

INTRODUCTION AND PREVIOUS WORK A. Networks of Workstations Networked workstations have been available for many years. Since a modern network of workstations represents a total computational capability on the order of a supercomputer, there is strong motivation to use a network of workstations (NOW) as a type of low-cost supercomputer. Note that a typical institution has access to a variety of computer resources that are network-inter- connected. These range from workstations to shared-memory multiprocessors to high-end parallel "supercomputers". In the context of this paper, a "network of workstations" or "NOW" is considered to be a network of heterogeneous computational resources, some of which may actually be worksta- tions. A run-time system for parallel programs on a NOW must have several important properties. First, scientific programmers needed a way to run supercomputer-class programs on such a system with little or no source-code modifications. The most common way to do this is to use an architecture-independent parallel programming standard to code the applications. This permits the same program source code to run on a NOW, a shared-memory multiprocessor, and the latest generations of parallel supercomputers, for example. Two major message-passing standards have emerged that can do this, PVM [1],[2] and MPI [3], as well as numerous distributed shared memory (DSM) systems Second, the ability to run large jobs on networked resources only proves attractive to workstation users if their individual workstations can still be used for more mundane tasks, such as word proces- This work was funded in part by NSF Grant No. EEC-8907070 Amendment 021 and by ONR Grant No. N00014-97-1-0116. r sing. Task migration is needed to "offload" work from a user's workstation and return the station to its "owner". This ability also permits dynamic load balancing and fault tolerance. (Note that, in this paper, a "task" is a piece of a parallel program, and a "program" or "job" is a complete parallel pro- gram. In the message-passing model used here, a program is always decomposed into multiple communicating tasks. A "platform" or "host" is a computer that can run jobs if idle, and can range from a workstation to an SMP.) Third, a run-time environment for NOW's needs the ability to track the availability and relative performance of resources as programs run, because it needs the information to conduct ongoing performance optimizations. This is true for several reasons. The relative speed of nodes in a network of workstations, even of homogeneous workstations, can vary widely. Workstation availability is a function of individual users, and users can create an external load that must be worked around. Programs themselves may run more efficiently if redistributed certain ways. Fourth, fault tolerance is extremely important in systems that involve dozens to hundreds of workstations. For example, if there is a 1% chance a single workstation will go down overnight, then there is only a (0.99) chance that a network of 75 stations will stay up overnight. A complete run-time system for NOW computing must therefore include an architecture-independent coding interface, task migration, automatic resource allocation, and fault-tolerance. It is also desirable for this system to support all of these features with no source-code modifications. These individual components already exist in various forms, and a review of them is in order. It is the goal of this work to combine these individual components into a single system. B. Parallel Programming Standards A wide variety of parallel and distributed architectures exist today to run parallel programs. Varying in their degree of interconnection, interconnection topology, bandwidth, latency, and scale of geographic distribution, they offer a wide range of performance and cost trade-offs and of applicability to solve different classes of problems. They can generally be characterized on the basis of their support for physical shared memory. Two major classes of parallel programming paradigms have emerged from the two major classes of parallel architectures. The shared-memory model has its origins in programming tightly coupled processors and offers relative convenience. The message-passing model is well suited for loosely coupled architectures and coarse-grained applications. Programming models that implement these paradigms can further be classified by their manifestation-either as extensions to existing programming languages or as custom languages. Because our system is intended for a network of workstations, it was felt that a message-pass- ing-based parallel programming paradigm more closely reflected the underlying physical structure. It was also felt that this paradigm should be expressed in existing scientific programming languages in order to draw on an existing base of scientific programmers. Both the PVM and MPI standards support these goals, and the MPI parallel programming environment was selected. Able to be called from C or FORTRAN programs, it provides a robust set of extensions that can send and receive messages between "tasks" working in parallel [3]. While a discussion of the relative merits of PVM and MPI is outside the scope of this paper, this decision was partially driven by ongoing work in MPI implementations that had already been conducted at Mississippi State and Argonne National Laboratory [4]. A more detailed discussion of the taxonomy of parallel paradigms and systems, and of the rationale for our decision, can be found in [5]. One desirable property of a run-time system is for its services to be offered "transparently" to applications programmers. Programs written to a common programming standard should not have to be modified to have access to more advanced run-time-system features. This level of transparency permits programs to maintain conformity to the programming standard and simplifies code development r C. Computing Systems Systems that can harness the aggregate performance of networked resources have been under development for quite some time. For example, one good review of cluster computing systems, conducted in 1995 and prepared at Syracuse University, listed 7 commercial and 13 research systems [6]. The results of the survey, along with a comparison to the Hector environment discussed in this paper, are summarized in [7]. It was found that the Hector environment had as many features as some of the other full-featured systems (such as Platform Computing's LSF, Florida State's DQS, IBM's Load Leveler, Genias' Codine, and the University of Wisconsin-Madison's Condor), and that its support for the simultaneous combination of programmer-transparent task migration and load-bal- ancing, fully automatic and dynamic load-balancing, support for MPI, and job suspension was unique. Additionally, the commitment to programmer transparency has led to the development of extensive run-time information-gathering about the programs as they run, and so the breadth and depth of the information that is gathered is unique. It should also be added that its support for typical commercial features, such as GUI and configuration management tools, was noticeably lacking, as Hector is a research project. It should also be mentioned at this point that there are (at least) two other research projects using the name "Hector" doing work in distributed computing and multiprocessing. The first is the well-known Hector multiprocessor project at the University of Toronto [8],[9]. The second is a system for supporting distributed objects in Python at the CRC for Distributed Systems Technology at the University of Queensland [10]. The Hector environment described in this paper is unrelated to either Three other research systems can allocate tasks across NOW's and have some degree of support for task migration. Figure 1 summarizes these systems. Condor/ CARMI Prospero MIST DQS Fully dynamic processor allocation and reallocation Only stops task under migration Y User-transparent load balancing Y User-transparent fault tolerance Y Y Y Works with MPI Y modifications to existing MPI/PVM program Uses existing operating system Y Y Y Y Figure 1: Comparison of Existing Task Allocators One such system is a special version of Condor [11], named CARMI [12]. CARMI can allocate PVM tasks across idle workstations and can migrate PVM tasks as workstations fail or become "non-idle". It has two limitations. First, it cannot claim newly available resources. For example, it does not attempt to move work onto workstations that have become idle. Second, it checkpoints all tasks when one task needs to migrate [13]. Stopping all tasks when only one task needs to migrate slows program execution since only the migrating task must be stopped. Another automated allocation environment is the Prospero Resource Manager, or PRM [14]. Each parallel program run under PRM has its own job manager. Custom-written for each program, the job manager acts like a "purchasing agent" and negotiates the acquisition of system resources as additional resources are needed. PRM is scheduled to use elements of Condor to support task migration and checkpointing and uses information gathered from the job and node managers to reallocate r resources. Notice that use of PRM requires a custom allocation "program" for each parallel applica- tion, and future versions may require modified operating systems and kernels. The MIST system is intended to integrate several development efforts and develop an automated task allocator [15]. Because it uses PRM to allocate tasks, the user must custom-build the allocation scheme for each program. MIST is built on top of PVM, and PVM's support of "indirect commu- nications" can potentially lead to administrative overhead, such as message forwarding, when a task has been migrated [16]. The implementation of MPI that Hector uses, with a globally available list of task hosts, does not incur this overhead. (Note that MPI does not have indirect communications, and so Hector does not have any overhead to support it.) Every task sends its messages directly to the receiving task, and the only overhead required after a task has migrated is to notify every other task of the new location. As will be shown below, this process has very little overhead, even for large parallel applications. The Distributed Queuing System, or DQS, is designed to manage jobs across multiple computers simultaneously [17]. It can support one or more queue masters which process user requests for resources on a first-come, first-served basis. Users prepare small batch files to describe the type of machine needed for particular applications. (For example, the application may require a certain amount of memory.) Thus resource allocation is performed as jobs are started. It currently has no built-in support for task migration or fault tolerance, but can issue commands to applications that can migrate and/or checkpoint themselves. It does support both PVM and MPI applications. The differences between these systems and Hector highlight two of the key differences among cluster computing systems. First, there is a trade-off between task migration mechanisms that are programmer-written versus those that are supported automatically by the environment. Second, there is a trade-off between centralized and decentralized decision-making and information-gathering D. Task Migration in the Context of Networked Resources Two strategies have emerged for creating the program's ``checkpoint'' or task migration image. First, checkpointing routines can be written by the application programmer. Presumably he or she is sufficiently familiar with the program to know when checkpoints can be created efficiently (for ex- ample, places where the running state is relatively small and free of local variables) and to know which variables are "needed" to create a complete checkpoint. Second, checkpointing routines can transfer the entire program state automatically onto another machine. The entire address space and registers are copied and carefully restored. By way of comparison, user-written checkpointing routines have some inherent "space" advan- tages, because the state's size is inherently minimized, and they may have cross-platform compatibility advantages, if the state is written in some architecture-independent format. User-written routines have two disadvantages. First, they add coding burden onto the programmer, as he or she must not only write but also maintain checkpointing routines. Second, checkpoints are only available at certain, predetermined places in the program. Thus the program cannot be checkpointed immediately on demand. It would appear from the Syracuse survey of systems [6] that most commercial systems only support user-written checkpointing for checkpointing and migration. One guesses that this is so because user-written checkpointing is much easier for resource management system developers-re- sponsibility for correct checkpointing is transferred to the applications programmer. As noted in earlier discussion, research systems such as PRM and DQS, also have user-written checkpointing, and at least one system (not discussed in the Syracuse report) has used this capability to perform cross-architecture task migration [18]. Condor and Hector use the complete-state-transfer method. This form of state transfer inherently only works across homogeneous platforms, because it involves actual replacement of a program's r state. It is also completely transparent to the programmer, requiring no modifications or additions to the source code. An alternate approach is to modify the compiler in order to retain necessary type information and pointer information. These two pieces of information (which data structures contain pointers and what the pointers "point to") are needed if migration is to be accomplished across heterogeneous platforms. At least one such system (the MSRM Library at Louisiana State University) has been implemented [19]. The MSRM approach may make automatic, cross-platform migration possible, at the expense of requiring custom compilers and of increased migration time. Once a system has a correct and consistent task migration capability, it is simple to add check- pointing. By having tasks transfer their state to disk (instead of to another task) it becomes possible to create a checkpoint of the program. This can be used for fault recovery and for job suspension. Thus both Hector and Condor provide checkpoint-and-rollback capability. However task migration is accomplished, there are two technical issues to deal with. First, the program's state must be transferred completely and correctly. Second, in the case of parallel pro- grams, any communications in progress must continue consistently. That is, the tasks must "agree" on the status of communications among themselves. Hector's solutions to these issues are discussed in section II.B. E. Automatic Resource Allocation There is a trade-off between a centralized allocation mechanism, in which all tasks and programs are scheduled centrally and the policy is centrally designed and administered, and a competitive, distributed model, in which programs compete and/or bid for needed resources and the bidding agents are written as part of the individual programs. Besides some of the classic trade-offs between centralized and distributed processing (such as overhead and scalability), there is an implied trade-off of the degree of support required by the applications programmer and of the intelligence with which programs can acquire the resources they need. The custom-written allocation approach places a larger burden on the applications programmer, but permits more well-informed acquisition of needed resources. Since the overall goal of Hector is to minimize programmer burden, it does not use any a priori information or any custom-written allocation policies. This is discussed further in section II.C. The degree to which a priori applications information can boost run-time performance has been explored for some time [20]. For example, Nguyen et al. have shown that extracting run-time information can be minimally intrusive and can substantially improve the performance of a parallel job scheduler [21]. Their approach used a combination of software and special-purpose hardware on a parallel computer to measure a program's speedup and efficiency and then used that information to improve program performance. However, Nguyen's work is only relevant for applications that can vary their own number of tasks in response to some optimization. Many parallel applications are launched with a specific number of tasks that does not vary as the program runs. Addition- ally, it requires the use of the KSR's special-purpose timing hardware. Gibbons proposed a simpler system to correlate run-times to different job queues [22]. Even this relatively coarse measurement was shown to improve scheduling, as it permits a scheduling model that more closely approaches the well-known Shortest Job First (SJF) algorithm. Systems can develop reasonably accurate estimates of a job's completion time based on historical traces of other programs submitted to the same queue. Since this information is coarse and gathered historically, it cannot be used to improve the performance of a single application at run-time. (It can, however, improve the efficiency of the scheduler that manages several jobs at once.) Some recent results by Feitelson and Weil have shown the surprising result that user estimates of run-time can make the performance of a job scheduler worse than the performance with no estimates at all [23]. While the authors concede that additional work is needed in the area, it does highlight that r user-supplied information can be unreliable, which is an additional reason why Hector does not use it. These approaches have shown the ability of detailed performance information to improve job scheduling. However, to summarize, these approaches have several shortcomings. First, some of them require special-purpose hardware. Second, some systems require user modifications to the applications program in order to keep track of relevant run-time performance information. Third, the information that is gathered is relatively coarse. Fourth, some systems require applications that can dynamically alter the number of tasks in use. Fifth, user-supplied information can be not only inaccurate but also misleading. The goal of Hector's resource allocation infrastructure is to overcome these shortcomings. There is another trade-off in degrees of support for dynamically changing workloads and computational resource availability. The ideal NOW distributed run-time system can automatically allocate jobs to available resources and move them around during the program run, both in order to maximize performance and in order to release workstations back to users. Current clustering systems support this goal to varying degrees. For example, some systems launch programs when enough resources (such as enough processors) become available. This is the approach taken by IBM's LoadLeveler, for example [6]. Other systems can migrate jobs as workstations become busy, such as Condor [11]. It appears that, as of the time of the Syracuse survey, only Hector attempts to move jobs onto newly idle resources as well. F. Goals and Objectives of Hector Because of the desire to design a system that supports fully transparent task migration, fully automatic and dynamic resource allocation, and transparent fault tolerance, the Hector distributed run-time environment is now being developed and tested at Mississippi State University. These requirements necessitated the development of a task-migration method and a modified MPI implementation that would continue correct communications during task migration. A run-time infrastructure that could gather and process run-time performance information was simultaneously created. The primary aim of this paper is to discuss these steps in more detail, as well as the steps needed to add support for fault tolerance. Hector is designed to use a central decision maker, called the "master allocator" or MA to perform the optimizations and make allocation decisions. It uses a small task running on each candidate processor to detect idle resources and monitor the performance of programs during execution. These tasks are called "slave allocators" or SA's. The amount of overhead associated with an SA is an important design consideration. An individual SA currently updates its statistics every 5 seconds. (This time interval is a compromise between timeliness and overhead.) This process takes about 5 ms on a Sun Sparcstation 5, and so corresponds to an extra CPU load of 0.1% [24]. The process of reading the task's CPU usage adds 581 #s per task every time that the SA updates its statistics (every 5 s). Adding the reading of detailed usage information therefore adds about 0.01% CPU load per task. For example, an SA supervising 5 MPI tasks will add a CPU load of 0.15%. When an MPI program is launched, individual pieces or "tasks" are allocated to available machines and migrated as needed. The SA's and the MA communicate to maintain awareness of the state of all running programs. The structure is diagrammed below in Figure 2. r Slave Allocator Master Allocator Performance Information System Info Commands Commands Performance Information Other Slave Allocators Figure 2: Structure of Hector Running MPI Programs Key design features and the design process are described below in section II and benchmarks and tests that measure Hector's performance are described in section III. This paper concludes with a discussion of future plans. II. GOALS, OBSTACLES, AND ACCOMPLISHMENTS A. Ease of Use A system must be easy to use if it is to gain widespread acceptance. In this context, ease of use can be supported two different ways. First, adherence to existing, widely accepted standards allows programmers to use the environment with a minimal amount of extra training. Second, the complexities of task allocation and migration and of fault tolerance should be "hidden" from unsophisticated scientific programmers. That is, scientific programmers should be able to write their programs and submit them to the resource management system without having to provide additional information about their program. 1. Using Existing Standards Hector runs on existing workstations and SMP's using existing operating systems, currently Sun systems running SunOS or Solaris and SGI systems running Irix. Several parts of the system, such as task migration and correctness of socket communications, would be simpler to support if modifications were made to the operating system. However, this would dramatically limit the usefulness of the system in using existing resources, and so the decision was made not to modify the operating system. The MPI and PVM standards provide architecture-independent parallel coding capability in both C and FORTRAN. MPI and PVM are supported on a wide and growing body of parallel architectures ranging from networks of workstations to high-end SMP's and parallel mainframes. these represent systems that have gained and are gaining widespread acceptance, there already exists a sizable body of programmers that can use it. Hector supports MPI as its coding standard. 2. Total Transparency of Task Allocation and Fault Tolerance Experience at the Mississippi State NSF Engineering Research Center indicates that most scientific programmers are unwilling (or unable) to provide such information as program "size", estimated rrun-time, or communications topology. This situation exists for two reasons. First, such programmers are solving a physical problem and so programming is a means to an end. Second, they may not have enough detailed knowledge about the internal workings of computers to provide information useful to computer engineers and scientists. Hector is therefore designed to operate with no a priori knowledge of the program to be executed. This considerably complicates the task allocation process, but is an almost-necessary step in order to promote "transparency" of task allocation to the programmer and, as a result, ease of use to the scientific programmer. Not currently supported, future versions of Hector may be able to benefit from user-supplied a priori information. A new implementation of MPI, named MPI-TM, has been created to support task migration [25] and fault tolerance. MPI-TM is based on the MPICH implementation of MPI [4]. In order to run with these features, a programmer merely has to re-link the application with the Hector-modified version of MPI. The modified MPI implementation and the Hector central decision-maker handle allocation and migration automatically. The programmer simply writes a "normal" MPI program and submits it to Hector for execution. Hector exists as the MPI-TM library and a collection of executables. The library is linked with applications and provides a self-migration facility, a complete MPI implementation, and an interface to the run-time system. The executables include the SA, MA, a text-based command-line interface to the MA, and a rudimentary Motif-based GUI. Its installation is roughly as complicated as installing a new MPI implementation and a complete applications package. 3. Support for Multiple Platforms Hector is supported on Sun computers running SunOS and Solaris and on SGI computers running Irix. The greatest obstacle under Solaris is its dynamic linker which, due to its ability to link at run-time, can create incompatible versions of the same executable file. This creates the undesirable situation that migration is impossible between nearly, but not completely, identical machines, and has the consequence of dividing the Sun computers into many, smaller clusters. This situation exists because of the combination of two factors. First, Hector performs automat- ic, programmer-transparent task migration without compiler modifications. Thus it cannot move pointers and must treat the program's state as an unalterable binary image. Second, dynamically linked programs may map system libraries and their associated data segments to different virtual addresses in runs of one program on different machines. The solution adopted by Condor is to re-write the linker (more accurately, to replace the Solaris dynamic linker with a custom-written one) to make migration of system library data segments possible [26]. This option is under consideration in Hector, but is not currently supported. B. Task Migration 1. Correct State Transfer The state of a running program, in a Unix environment, can be considered in six parts. First, the actual program text may be dynamically linked, and has references to data that may be statically or dynamically located. Second, the program's static data is divided into initialized and uninitialized sections. Third, any use of dynamically allocated data is stored in "the heap". Fourth, the program's stack grows as subroutines and functions are called, and is used for locally visible data and dynamic data storage. Fifth, the CPU contains internal registers, usually used to hold results of intermediate calculations. Sixth, the Unix kernel maintains some system-level information about the program, such as file descriptors. This is summarized below in Figure 3. r CPU User Memory Kernel Memory Visible to user Not visible to user Text Static Data Heap Registers Stack Kernel Structs Wrapper functions keep track of kernel information in a place visible to the user. Figure 3: State of a Program During Execution The first five parts of the state can, in principle, be transferred between two communicating user-level programs. One exception occurs when programs are dynamically linked, as parts of the program text and data may not reside at the same virtual address in two different instantiations of the same program. As discussed above, this matter is under investigation. The sixth part of a program's state, kernel-related information, is more difficult to transfer because it is "invisible" to a user-level program. This information may include file descriptors and pointers, signal handlers, and memory-mapped files. Without kernel source code, it is almost impossible to read these structures directly. If the operating system is unmodified, the solution is to create "wrapper" functions that let the program keep track of its own kernel-related structures. All user code that modifies kernel structures must pass through a "trap" interface. (Traps are the only way user-level code can execute supervisor-level functions.) The Unix SYSCALL.H file documents all of the system calls that use traps, and all other system calls are built on top of them. One can create a function with the same name and arguments as a system call, such as open(). The arguments to the function are passed into an assembly-language routine that calls the system trap proper- ly. The remainder of the function keeps track of the file descriptor, path name, permissions, and other such information. The lseek() function can keep track of the location of the file pointer. Calls that change the file pointer (such as read() and write() ) also call the instrumented lseek(), so that file pointer information is updated automatically. This permits migrated tasks to resume reading and writing files at the proper place. It was discovered that the MPI environment for which task migration was being added [5] also uses signals and memory mapping. (The latter is due to the fact that gethostbyname() makes a call to mmap.) All system calls that affect signal handling and memory mapping are replaced with wrapper functions as well. The task migration system requires knowledge of a running program's image in a particular operating system, the development of a small amount of assembly language, and reliance on certain properties pertaining to signal-handling, and these all affect the portability of the task migration sys- tem. The assembly language is needed because this is the only way to save and restore registers and call traps. Since the task migration routine is inside a signal handler, it is also necessary for the re-started program to be able to exit the signal handler coherently. Other systems that perform a similar style of system-supported task migration, such as MIST and Condor, have also been ported to Linux, Alpha, and HP environments [27],[15]. This seems to indicate that this style of task migration is reasonably portable among Unix-based operating sys- tems, probably because these different operating systems have strong structural similarities. It is r interesting to note that no system-level migration support for Windows/NT-based systems has been reported. The exact sequence of steps involved in the actual state transfer are described in more detail in [25]. Two tests confirm this method's speed and stability, and are described below in section III. 2. Keeping MPI Intact: A Task Migration Protocol The state restoration process described above is not guaranteed to preserve communications on sockets. This is because at any point in the execution of a program, fragments of messages may reside in the kernel's buffers on either the sending or receiving side. The solution is to notify all tasks when a single task is about to migrate. Each task that is communicating with the task under migration sends an "end-of-channel" message to the migrating task and then closes the socket that connects them. The tasks then mark the migrating task as "under migration" in its table of tasks, and attempts to initiate communications will block until migration is complete. Once the task under migration receives all of its "end-of-channel" messages, it can be assured that no messages are trapped in the buffers. That is, it knows that all messages reside in its data segment, and so it can be migrated "safely". Once the state has been transferred, another global update is needed so that other tasks know its new location and know that communications can be resumed with it. Tasks that are not migrating remain able to initiate connections and communicate with one another. The MPI-1.1 standard (the "original" MPI standard) only permits static task tables. That is, the number of tasks used by a parallel program is fixed when the program is launched. (It is important to note that the static number of tasks in a program is an MPI-1.1 limit, not a limitation of Hector. This is also one of the important differences between PVM and MPI-1.1.) Thus updates to this table do not require synchronization with MPI and do not "confuse" an MPI program. The MPI-2 standard (a newer standard currently in development) permits dynamically changing task tables, but, with proper use of critical sections, task migration will not interfere with programs written under the MPI-2 standard. A series of steps is needed to update the task table globally and atomically. Hector's MA and SA's are used to provide synchronization and machine-to-machine communications during migration and task termination. The exact sequence of steps required to synchronize tasks and update the communication status consistently is described in detail in [25]. It should be noted that if the MA crashes in the middle of a migration, the program will deadlock, because the MA is used for global synchronization and to guarantee inter-task consistency. 3. Task Termination Protocol Task termination presents another complication. If a task is migrating while or after another task terminates, the task under migration never receives an "end-of-channel" message from the terminated task. Two measures are taken to provide correct program behavior. First, the MA limits each MPI program to only one migration or only one termination at a time. It can do this because of the handshaking needed both to migrate and to terminate. Second, a protocol involving the SA's and MA's was developed to govern task termination and is described below. 1. A task preparing to terminate notifies its SA. The task can receive and process table updates and requests for end-of-channel (EOC) messages, but will block requests to migrate. It cannot be allowed to migrate so that the MA can send it a termination signal. 2. The SA notifies the MA that the task is ready to terminate. 3. Once all pending migrations and terminations have finished, the MA notifies the SA that the task "has permission" to terminate. It will then block (and enqueue) further termination and migration requests until this termination has ended. 4. The SA notifies the task. 5. The task sends the SA a final message before exiting. 6. The SA notifies the MA that the task is exiting, and so the MA can permit other migrations and terminations. r Notice that an improperly written program may attempt to communicate with a task after the task has ended. In the world of message-passing-based parallel programming, this is a programmer's mistake. Behavior of the program is undefined at this point, and the program itself will deadlock under Hector. (The program deadlocks, not Hector.) 4. Minimizing Migration Time The operating system already has one mechanism for storing a program's state. A core dump creates a file that has a program's registers, data segment, and stack. The first version of state transfer used this capability to move programs around. There are two advantages to this approach. First, it is built into the operating system. Second, there are symbolic debuggers and other tools that can extract useful information from core files. There are some disadvantages to this approach. First, multiple network transfers are needed if the disk space is shared over a network. This means that the state is actually copied multiple times. Second, the speed of transfer is limited further by the speed of the disk and by other, unrelated programs sharing that disk. One way around all of these shortcomings is to transfer the state directly over the network. Originally implemented by the MIST team [15], network state transfer overcomes these disadvantages. The information is written over the network in slightly modified core-file format. (The only modification is that unused stack space is not transmitted. There is no other penalty for using the SunOS core file format.) The information is written over a network socket connection by the application itself, instead of being written to a file by the operating system. Notice that this retains the advantage of core-file tool compatibility. Experiments show that it is over three times faster [25], as will be shown below. C. Automatic Resource Allocation 1. Sources of Information Hector's overall goal is to attempt to minimize programmer overhead. In the context of awareness of program behavior, this incurs the expense of not having access to potentially beneficial program-specific information. This approach was used based on experiences with scientific programmers within the authors' research center, who are unwilling to invest time and effort to use new systems because of the perceived burden of source-code modifications. This approach dictates that Hector be able to operate with no a priori applications knowledge, which, in turn, increases the requirement for the depth and breadth of information that is gathered at run-time. This lack of a priori information makes Hector's allocation decision-making more difficult. However, experiments confirm that the information it is able to extract at run-time can improve performance, and its ability to exploit newly idle resources is especially helpful. 2. Structure There is a range of ways that resource allocation can be structured, from completely centrally located to completely distributed. Hector's resource allocation uses features of both. The decision-making portion (and global synchronization) resides in the MA and is therefore completely central- ized. The advantage of a single, central allocation decision-maker is that it is easier to modify and test different allocation strategies. Since the UNIX operating system will not permit signals to be sent between hosts, it is necessary to have an executive process running on each candidate host. Since it is necessary to have such pro- cesses, they can be used to gather performance information as well. Thus its information-gathering and information-execution portions are fully distributed, being carried out by the SA's. 3. Collecting Information There are two types of information that the master allocator needs in order to make decisions. First, it needs to know about the resources that are potentially available, such as which hosts to consider rand how powerful they are. Second, it needs to know how efficiently and to what extent these resources are being used, such as how much external (non-Hector) load there is and how much load the various MPI programs under its control are imposing. The relative performance of each candidate host is determined by the slave allocator when it is started. (It does so by running the Livermore Loops [28], which actually measure floating point per- formance.) It is also possible for the slave allocator to measure disk availability and physical memory size, for example. This information is transmitted to the master allocator, which maintains its own centralized database of this information. Current resource usage is monitored by analyzing information from the kernel of each candidate processor. Allocation algorithms draw on idle time information, CPU time information, and the percentage of CPU time devoted to non-Hector-related tasks. The percentage of CPU time is used to detect "external workload", such as an interactive user logging in, which is a criterion for automatic migration. The Hector MPI library contains additional, detailed self-instrumentation that logs the amount of computation and communication time each task expends. This data is gathered by the SA's by using shared-memory and is forwarded to the MA. A more detailed discussion of this agent-based approach to information-gathering, as well as testing and results, may be found in [29]. 4. Making Decisions One of the primary advantages of this performance-monitoring-based approach is its ability to claim idle resources rapidly. As will be shown below, tests on busy workstations during the day show that migrating to newly available resources can reduce run time and promote more effective use of workstations. Further implementation and testing of more sophisticated allocation policies are also under way. D. Fault Tolerance The ability to migrate tasks in mid-execution can be used to suspend tasks. In fact, fault tolerance has historically been one major motivation for task migration. In effect, each task transfers its state into a file to "checkpoint" the program. When a node failure has been detected, the files can be used to "roll back" the program to the state of the last checkpoint. While all calculations between the checkpoint and node failure are lost, the calculations up to the checkpoint are not, which may represent a substantial time savings. Also, known unrelated failures and/or routine maintenance may occur or be needed in the middle of a large program run, and so the ability to suspend tasks is helpful. It can be shown that in order to guarantee program correctness, all tasks must be checkpointed consistently[13]. That is, the tasks must be at a consistent point in their execution and in their message exchange status. For example, all messages in transit must be fully received and transmission of any new messages must be suspended. As was the case with migration and termination, a series of steps are needed to checkpoint and to roll back parallel programs. 1. Checkpointing Protocol The following steps are taken to checkpoint a program. 1. The MA decides to checkpoint a program for whatever reason. (This is currently supported as a manual user command, and may eventually be done on a periodic basis.) It waits until all pending migrations and terminations have finished, and then it notifies all tasks in the program, via the SA's, to prepare for checkpointing. 2. The tasks send end-of-channel (EOC) messages to all connected tasks, and then receive EOC's from all connected tasks. Again, this guarantees that there are no messages in transit. 3. Once all EOC's have been exchanged, the task notifies its SA that it is ready for checkpointing and informs the SA of the size of its state information. This information is passed on to the MA. 4. Once the MA has received confirmation from every task, it is ready to begin the actual check-pointing process. It notifies each task when the task is to begin checkpointing. 5. After each task finishes transmitting its state (or writing a file), it notifies the MA. Note that it is possible for more than one task to checkpoint at a time, and experiments with the ordering of checkpointing are described below. 6. After all tasks have checkpointed, the MA writes out a small "bookkeeping" file which contains state information pertinent to the MA and SA's. (For example, it contains the execution time to the point of checkpointing, so that the total execution time will be accurate if the job is rolled back to that checkpoint.) 7. The MA broadcasts either a "Resume" or "Suspend" command to all tasks. The tasks either resume execution or stop, respectively. The former is used to create a "backup copy" of a task in the event of future node failure. The latter is used if it is necessary to remove a job temporarily from the system. 2. Rollback Protocol The following steps are taken to roll back a checkpointed program. 1. The MA is given the name of a "checkpoint file" that provides all necessary information to restart the program. 2. It allocates tasks on available workstations, just as if the program were being launched. 3. Based on its allocation, it notifies the SA on the first machine. 4. The SA restarts the task from the "state file", the name of which is found in the checkpoint file and sent to the SA. 5. The task notifies its SA that it restarted properly and waits for a "table update". 6. Once the MA receives confirmation of one task's successful restart, it notifies the SA of the next task. It continues to do this until all tasks have restarted. Task rollback is sequential primarily for performance reasons. The file server that is reading the checkpoints and sending them over sockets to the newly launched tasks will perform more efficiently if only one checkpoint is sent at a time. 7. As confirmation arrives at the MA, it builds a table similar to that used by the MPI tasks themselves. It lists the hostnames and Unix PID's of all the tasks in the parallel program. Once all tasks have restarted, this table is broadcast to all tasks. Note that this broadcast occurs via the Hector run-time infrastructure and is "invisible" to the MPI program. It does not use an MPI broadcast, as MPI is inactive during rollback. 8. Each task resumes normal execution once it receives its table update, and so the entire program is restarted. 3. The Checkpoint Server As is the case with task migration, there are two ways to save a program's state. One way is for each program to write directly to a checkpoint file. The other way is to launch a "checkpoint server" on a machine with a large amount of physically mounted disk space. (The latter concept was first implemented by the Condor group [30].) Each task transmits its state via the network directly to the server, and the server writes the state directly to its local disk. The reason each task cannot write its state to its local disk is obvious-if the machine crashes, the "backup copy" of the state would be lost as well. The checkpoint server method is expected to be faster, because it uses direct socket connections and local disk writes, which are more efficient than writing files over a network. Note that many of the local disk caching strategies used by systems like NFS do not work well for checkpoints, because checkpoint files are typically written once and not read back [30]. Different, novel scheduling strategies for checkpoint service are described and tested below. 4. Other Issues The MA is not fault-tolerant. That is, the MA represents a single point of failure. The SA's have been modified to terminate themselves, and the tasks running under them, if they lose contact with r the MA. (This feature was intentionally added because total termination of programs distributed across dozens of workstations can be quite tedious unless it is automated.) If this feature is disabled, then SA's and their tasks could continue working without the MA, although all task migrations and job launches would cease, and job termination would deadlock. The checkpoints collected by the system will enable a job to be restarted after the MA and SA's have been restarted, and so the check- point-and-rollback-based fault tolerance can tolerate a fault in the run-time infrastructure. Another approach to solve this problem, and to support more rapid fault tolerance, would be to incorporate an existing group communication library (such as Isis or Horus [31]) and use its mes- sage-duplication facility. One possible design is described in [32]. Means of rapid fault detection can also be added to future versions of Hector [32]. Each SA sends performance information to the MA periodically. (The current period is 5 seconds, which may grow as larger tests are performed.) If the SA does not send a performance update after some suitable timeout, it can be assumed that the node is not running properly, and all jobs on that node can be rolled back. This strategy will detect heavily overloaded nodes as well as catastrophically failed nodes. III. BENCHMARKS AND TESTS A. Ease of Use As an example, an existing computational fluid dynamics simulation was obtained, because it was large and complex, having a total data size of around 1 GByte and about 13,000 lines of Fortran source code. The simulation had already been parallelized, coded in MPI, and tested on a parallel computer for correctness. With no modifications to the source code, the program was relinked and run completely and correctly under Hector. This highlights its ability to run "real", existing MPI programs. B. Task Migration Two different task migration mechanisms were proposed, implemented, and tested. The first used "core dump" to write out a program's state for transfer to a different machine. The second transferred the information directly over a socket connection. In order to compare their relative speeds, tasks of different sizes were migrated times each between two Sparcstations 10's connected ethernet. The tests were run during normal daily operations at the Engineering Research Center [25]. The results are show below in Figure 4. Program Size (kbytes)515 Time toMigrate (sec) Core File Transfer Network Transfer Figure 4: Time to migrate tasks of different sizes r Consider, for example, the program of size 2.4 Mbytes. The "core dump" version migrated in about 12.6 to 13.1 seconds. This is an information transfer rate of roughly 1.5 MBit/sec. The "net- work state transfer" version migrated in 4.1 seconds, or an information rate of about 4.6 MBit/sec. The latter is close to the practical limit of ordinary 10 Mbit ethernet under typical conditions, which shows that task transfer is limited by network bandwidth. Overall, network state transfer was about 3.2 times faster. The same test program was run with minimal size and with network state transfer. The task migrated over 400,000 times without error, and only stopped when a file server went down due to an unrelated fault. (The file server crash had the effect of locking up the local machine, and so Hector became locked up as well, as the fork-and-exec needed to launch a new task could not be executed.) The overhead associated with broadcasting table updates was measured by migrating tasks inside increasingly large parallel programs. This test was run between a Sparcstation 5 and an ether- net-connected Sparcstation 10 on separate subnets, and was run under normal daily operating conditions at the Engineering Research Center. The program did the following. First, one task establishes communications with some number of other tasks. Second, that task is migrated. As explained above, this entails sending notification, receiving end-of-channel messages, and transferring the state of the program. Third, the newly migrated task re-establishes communications with all other tasks. The time from ordering the task to migrate to re-establishing all connections was measured and the test was repeated 50 times. The total number of tasks in the program ranged from 10 to 50, and the results are shown below in Figure 5. The crosses are actual data points, and the line is a linear regression best-fit through the data. A linear regression was chosen because the time to migrate is a linear function of both the number of other tasks and the size of the migration image.135 Number of Other Tasks Time toMigrate (sec) Figure 5: Time to migrate a task as the number of other tasks increases The slope of the best-fit line corresponds to the incremental amount of time required to notify one other task of impending migration, wait for the EOC message, and then notify the other task that migration was completed. Since the program size changed only negligibly as the number of tasks increased, the time to migrate was effectively constant. The results are show in Figure 5, and the slope corresponds to about 75 ms per task per migration. That is, the process of notifying, receiving end-of-channel, and re-establishing communications takes, on the average, about 75 ms over 10 Mbit/s ethernet. The Y-intercept is roughly the time it took the task to migrate, and is about 184 ms. If the transfer took place at 10 Mbit/s, this corresponds to a migration image size of about 230 Kbytes. C. Automatic Resource Allocation In one early test, an MPI program was run under two different scenarios to determine the performance benefits of "optimal" scheduling over "first come, first served" scheduling [5]. A three-process matrix multiply program was used, and a 1000 x 1000 matrix was supplied as the test case. The machines used in the test ranged from a 40 MHz SPARC-based system to a 4-processor 70 MHz Hyper-SPARC-based system. All were free of external loads, except as noted below. (The matrix multiply is a commonly used test program because its data size is quite scalable and the program's structure is simple and easily modified.) First, tests were run to validate the advantages of using a processor's load as a criterion for task migration. The third workstation was loaded with an external workload. The runtime dropped from 1316 seconds to 760 seconds when the third task was migrated onto an idle workstation. Second, tests were run to show the advantages of migrating tasks in the event faster workstations come available. When a faster machine became available two minutes into the computation, the run-time dropped from 1182 seconds to 724 seconds. Algorithms used by other task allocation systems would not have migrated the slow tasks once the program was launched. More extensive tests were run once the master allocator was more completely developed. The same matrix multiply program was used, and the allocation policy was switched between two varia- tions. The first variation only migrated tasks when workstations became busy, usually from exter- nal, interactive loads. The second variation considered migrations whenever a workstation became idle. A test system was written that switched the MA between these two allocation policies on alternate program runs, and a total of 108 runs were made during several days at the research center on 3 fairly well-used Sparcstation 2's. When the "only-migrate-when-busy" policy was used, the run times varied from 164 to 1124 seconds. When the "also-migrate-when-idle" policy was used, the run times varied from 154 to 282 seconds. In the latter case, the program averaged 0.85 task migrations per run to idle worksta- tions. While migration to idle workstations had little effect on minimum run times, they dramatically reduced the maximum run time. The average run time dropped from 359 seconds to 177 seconds due to this reduction in maximum run time. The distribution of run times is shown below in Figure 6. Run Times - 1st Policy Number of Trials Run Times - 2nd Policy Figure Distribution of Run Times for Two Task Migration Policies This improvement in run time was noticed in cases where relatively small programs were being migrated. As programs increase in size, the penalty for migration increases. For example, runs of a fairly large electromagnetics simulation showed a noticeable increase in run time as the number of migrations increased [7]. r One point of interest is that interactive were sometimes unaware that their machines were also being used to run jobs under Hector (including this paper's primary author). Most users are accustomed to a small wait for applications such as word processors to page in from virtual memory. Since the time to migrate tasks off of their machines was on the order of the time required to page in pro- grams, the additional interactive delay was not noticed. It should be noted that this was only the case for small test programs, as truly rapid migration is only practical when the program state fits within physical memory. As was shown in tests with larger programs, once the program size exceeds physical memory the process of migration causes thrashing [7]. More extensive testing, showing runs of a variety of applications on both Sun and SGI workstation clusters and runs alongside actual student use can be found in [7]. D. Fault Tolerance A series of tests were run to evaluate the relative performance of different means of creating and scheduling checkpoints. The same matrix multiply program was used to test different combinations of strategies. The matrices were 2 100X100, 500X500, and 1000X1000 matrices for small, medium, and large checkpoint sizes, respectively. The total checkpoint sizes were 2 Mbytes, 11 Mbytes, and Mbytes respectively. The following tests were also run during the day and under normal network usage conditions. First, two strategies were tested that involved writing directly to a checkpoint file using NFS. All tasks were commanded to checkpoint simultaneously ("All-at-Once") or each task was commanded to checkpoint by itself ("One-at-a-Time"). Second, three server-based strategies were evaluated for overall effectiveness. First, the server was notified of a single task to checkpoint, and so tasks were checkpointed one at a time. Second, the server was notified of all tasks and their sizes, but only one task at a time was permitted to checkpoint. (This is labelled "combined" in the figures below.) Third, the server was notified of all tasks, and then all tasks were ordered to checkpoint simultaneously. The individual tasks were instrumented to report the amount of time required to send the data either to a file (over NFS) or to a checkpoint server. The size of the program state and the time to checkpoint the state were used to calculate bandwidth usage, and are shown below for both cases- without (Figure 7) and with (Figure 8) a checkpoint server. Strategy Small Medium Large All-at-Once 2.21 1.46 1.45 One-at-a-Time 4.83 5.12 5.55 Figure 7: Bandwidth Usage of Different Strategies with No Server (MBit/sec) Strategy Small Medium Large All-at-Once 2.37 1.55 1.71 Combined 6.57 7.15 7.25 One-at-a-Time 7.26 7.47 7.20 Figure 8: Bandwidth Usage of Different Strategies with a Server (MBit/sec) The results are easy to understand. First, using NFS instead of a server involves some bandwidth penalty. Using a server permits data to be transferred without additional protocol overheads. Se- cond, the all-at-once strategy clogs the ethernet and substantially reduces its performance. The combined and one-at-a-time strategies have essentially identical performance from the point of view of the parallel program because each task sends its data to the server without interference from other tasks. r The server works by forking a separate process for each task to be checkpointed. The process "mallocs" space for the entire state, reads the state over a socket, and writes the state to disk. (It knows the size of the program state because the task notifies its SA of the size when it indicates readiness to checkpoint in step 3 of the checkpoint protocol. This information is passed, via the MA, to the checkpoint server.) Data were captured by the checkpoint server itself, and the results are shown below in Figure 9. The individual points are some of the actual data points, and the lines are linear regression curves fit through the actual data. Each policy was tested times over the range of sizes shown on the graph.200 10000 20000 30000200 10000 20000 30000 Total Size of Program Checkpoint (Kbytes) Time to Checkppoint Program (sec) One-at-a-Time All-at-Once Combined Figure 9: Time to Checkpoint Tasks under Different Checkpoint Server Policies The reciprocal of the slope of each line corresponds to the bandwidth obtained by that strategy. They are summarized in Figure 10. Bandwidth (MBit/sec) One-at-a-Time 5.43 Combined 6.86 All-at-Once 4.87 Figure 10: Impact of Different Server Strategies The combined strategy makes best use of available bandwidth, and the "all-at-once" strategy has the worst performance. This is typical for ethernet-based communications. In fact, the penalty shown here is less than that expected by the authors. The conclusion is that other traffic on the net-work rat the time of program runs already has a significant impact, and that using the "all-at-once" strategy compounds the problem. Note that the time to checkpoint, as measured by the server and shown in Figure 9, includes the time to fork processes and malloc space, the time to read the data from the network, and the time to write it to disk. The "bandwidth" is actually a function of all of these times. This is the reason why the combined strategy shows improved performance-the process of forking new processes and making space in memory for the state images is done in parallel. Also notice that when the time required to write to disk and fork tasks is added, the bandwidth of the combined strategy only drops slightly, which indicates that the overhead of the fork-and-malloc process is small. IV. FUTURE WORK AND CONCLUSIONS A. Testing of Dynamic Load-Balancing Some initial concepts for dynamic load balancing have been implemented. These are undergoing testing, and will probably be modified as testing proceeds. Since only one function is needed to optimize tasks, it is very simple to provide alternate optimization strategies for testing. B. Job Priority Currently all jobs run at the same priority level. Some concept of "priority" needs to be added in future versions to permit job scheduling more like that found in industry. The combination of priority-based decision making and built-in fault tolerance will be explored as the basis of a fault-toler- ant, highly reliable computing system. C. Conclusions Running parallel programs efficiently on networked computer resources requires architecture-independent coding standards, task migration, resource awareness and allocation, and fault toler- ance. By modifying an MPI implementation and developing task migration technology, and by developing a complete run-time infrastructure to gather performance information and make optimiza- tions, it became possible to offer these services transparently to unmodified MPI programs. The resulting system features automatic and fully dynamic load-balancing, task migration, the ability to release and reclaim resources on demand, and checkpoint-and-rollback-based fault tolerance. It has been demonstrated to work with large, complex applications on Sun and SGI workstations and SMP's. Tests confirm that its ability to reclaim idle resources rapidly can have beneficial effects on program performance. A novel checkpoint server protocol was also developed and tested, making fault tolerance more efficient. Since it combines and supports, to varying degree, all of the features needed for NOW parallel computing, it forms the basis for additional work in distributed computing. V. --R "Visualization and debugging in a heterogeneous environment," PVM: Parallel Virtual Machine Cambridge Mass: The MIT Press "Monitors, Message, and Cluster: The p4 Parallel Program System," "Hector: Automated Task Allocation for MPI" "Cluster Computing Review" "Using Hector to Run MPI Programs over Networked Workstations" "Experiences with the Hector Multiprocessor" "Experiences with the Hector Multiprocessor" "Hector: Distributed Objects in Python" "The Condor Distributed Processing System" "Providing Resource Management Services to Parallel Applica- tions" "Consistent Checkpoints of PVM Applications" "The Prospero Resource Manager: A Scalable Framework for Processor Allocation in Distributed Systems" "MIST: PVM with Transparent Migration and Checkpointing" "MPVM: A Migration Transparent Version of PVM" DQS User Manual - DQS Version 3.1 "Portable Checkpointing and Recovery" "Memory Space Representation for Heterogeneous Network Process Migration" "Theory and Practice in Parallel Job Scheduling" "Using Runtime Measured Workload Characteristics in Parallel Processing Scheduling" "A Historical Application Profiler for Use by Parallel Schedulers" "Utilization and Predictability in Scheduling the IBM SP2 with Backfilling" "User-Transparent Run-Time Performance Optimization" "A Task Migration Implementation for the Message-Passing Interface" "Checkpoint and Migration of UNIX Processes in the Condor Distributed Processing System" Page, http://www. The Livermore Fortran Kernels: A Computer Test Of The Numerical Performance Range "An Agent-Based Architecture for Dynamic Resource Management" "Managing Checkpoints for Parallel Programs" "Software-based Replication for Fault Tolerance" "An Architecture for Rapid Distributed Fault Tolerance" --TR --CTR Samuel H. Russ , Katina Reece , Jonathan Robinson , Brad Meyers , Rajesh Rajan , Laxman Rajagopalan , Chun-Heong Tan, Hector: An Agent-Based Architecture for Dynamic Resource Management, IEEE Concurrency, v.7 n.2, p.47-55, April 1999 Angela C. Sodan , Lin Han, ATOP-space and time adaptation for parallel and grid applications via flexible data partitioning, Proceedings of the 3rd workshop on Adaptive and reflective middleware, p.268-276, October 19-19, 2004, Toronto, Ontario, Canada Hyungsoo Jung , Dongin Shin , Hyuck Han , Jai W. Kim , Heon Y. Yeom , Jongsuk Lee, Design and Implementation of Multiple Fault-Tolerant MPI over Myrinet (M^3), Proceedings of the 2005 ACM/IEEE conference on Supercomputing, p.32, November 12-18, 2005 Kyung Dong Ryu , Jeffrey K. Hollingsworth, Exploiting Fine-Grained Idle Periods in Networks of Workstations, IEEE Transactions on Parallel and Distributed Systems, v.11 n.7, p.683-698, July 2000

task migration;load balancing;fault tolerance;resource allocation;parallel computing

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A Fault-Tolerant Dynamic Scheduling Algorithm for Multiprocessor Real-Time Systems and Its Analysis.

AbstractMany time-critical applications require dynamic scheduling with predictable performance. Tasks corresponding to these applications have deadlines to be met despite the presence of faults. In this paper, we propose an algorithm to dynamically schedule arriving real-time tasks with resource and fault-tolerant requirements on to multiprocessor systems. The tasks are assumed to be nonpreemptable and each task has two copies (versions) which are mutually excluded in space, as well as in time in the schedule, to handle permanent processor failures and to obtain better performance, respectively. Our algorithm can tolerate more than one fault at a time, and employs performance improving techniques such as 1) distance concept which decides the relative position of the two copies of a task in the task queue, 2) flexible backup overloading, which introduces a trade-off between degree of fault tolerance and performance, and resource reclaiming, which reclaims resources both from deallocated backups and early completing tasks. We quantify, through simulation studies, the effectiveness of each of these techniques in improving the guarantee ratio, which is defined as the percentage of total tasks, arrived in the system, whose deadlines are met. Also, we compare through simulation studies the performance our algorithm with a best known algorithm for the problem, and show analytically the importance of distance parameter in fault-tolerant dynamic scheduling in multiprocessor real-time systems.

Introduction Real-time systems are defined as those systems in which the correctness of the system depends not only on the logical result of computation, but also on the time at which the results are produced [22]. Real- This work was supported by the Indian National Science Academy, and the Department of Science and Technology. time systems are broadly classified into three categories, namely, (i) hard real-time systems, in which the consequences of not executing a task before its deadline may be catastrophic, (ii) firm real-time systems, in which the result produced by the corresponding task ceases to be useful as soon as the deadline expires, but the consequences of not meeting the deadline are not very severe, and (iii) soft real-time systems, in which the utility of results produced by a task with a soft deadline decreases over time after the deadline expires [25]. Examples of hard real-time systems are avionic control and nuclear plant control. Online transaction processing applications such as airline reservation and banking are examples for firm real-time systems, and telephone switching system and image processing applications are examples for soft real-time systems. The problem of scheduling of real-time tasks in multiprocessor systems is to determine when and on which processor a given task executes [22, 25]. This can be done either statically or dynamically. In static algorithms, the assignment of tasks to processors and the time at which the tasks start execution are determined a priori. Static algorithms are often used to schedule periodic tasks with hard deadlines. However, this approach is not applicable to aperiodic tasks whose characteristics are not known a priori. Scheduling such tasks require a dynamic scheduling algorithm. In dynamic scheduling, when a new set of tasks (which correspond to a plan) arrive at the system, the scheduler dynamically determines the feasibility of scheduling these new tasks without jeopardizing the guarantees that have been provided for the previously scheduled tasks. A plan is typically a set of actions that has to be either done fully or not at all. Each action could correspond to a task and these tasks may have resource requirements, and possibly may have precedence constraints. Thus, for predictable executions, schedulability analysis must be done before a task's execution is begun. For schedulability analysis, tasks' worst case computation times must be taken into account. A feasible schedule is generated if the timing constraints, and resource and fault-tolerant requirements of all the tasks in the new set can be satisfied, i.e., if the schedulability analysis is successful. If a feasible schedule cannot be found, the new set of tasks (plan) is rejected and the previous schedule remains intact. In case of a plan getting rejected, the application might invoke an exception task, which must be run, depending on the nature of the plan. This planning allows admission control and results in reservation-based system. Tasks are dispatched according to this feasible schedule. Such a type of scheduling approach is called dynamic planning based scheduling [22], and Spring kernel [27] is an example for this. In this paper, we use dynamic planning based scheduling approach for scheduling of tasks with hard deadlines. The demand for more and more complex real-time applications, which require high computational needs with timing constraints and fault-tolerant requirements, have led to the choice of multiprocessor systems as a natural candidate for supporting such real-time applications, due to their potential for high performance and reliability. Due to the critical nature of the tasks in a hard real-time system, it is essential that every task admitted in the system completes its execution even in the presence of failures. Therefore, fault-tolerance is an important issue in such systems. In real-time multiprocessor systems, fault-tolerance can be provided by scheduling multiple versions of tasks on different processors. Four different models (techniques) have evolved for fault-tolerant scheduling of real-time tasks, namely, (i) Triple Modular Redundancy (TMR) model [12, 25], (ii) Primary Backup (PB) model [3], (iii) Imprecise Computational (IC) model [11], and (iv) (m; k)-firm deadline model [23]. In the TMR approach, three versions of a task are executed concurrently and the results of these versions are voted on. In the PB approach, two versions are executed serially on two different processors, and an acceptance test is used to check the result. The backup version is executed (after undoing the effects of primary version) only if the output of the primary version fails the acceptance test, either due to processor failure or due to software failure. In the IC model, a task is divided into mandatory and optional parts. The mandatory part must be completed before the task's deadline for acceptable quality of result. The optional part refines the result. The characteristics of some real-time tasks can be better characterised by (m; k)-firm deadlines in which m out of any k consecutive tasks must meet their deadlines. The IC model and (m; k)-firm task model provide scheduling flexibility by trading off result quality to meet task deadlines. Applications such as automatic flight control and industrial process control require dynamic scheduling with fault-tolerant requirements. In a flight control system, the controllers often activate tasks depending on what appears on their monitor. Similarly, in an industrial control system, the robot which monitors and controls various processes may have to perform path planning dynamically which results in activation of aperiodic tasks. Another example, taken from [3], is a system which monitors the condition of several patients in the intensive care unit (ICU) of a hospital. The arrival of patients to the ICU is dynamic. When a new patient (plan) arrives, the system performs admission test to determine whether the new patient (plan) can be admitted or not. If not, alternate action like employing a nurse can be carried out. The life criticality of such an application demands that the desired action to be performed even in the presence of faults. In this paper, we address the scheduling of dynamically arriving real-time tasks with PB fault-tolerant requirements on to a set of processors and resources in such a way that the versions of the tasks are feasible in the schedule. The objective of any dynamic real-time scheduling algorithm is to improve the guarantee ratio [24] which is defined as the percentage of tasks, arrived in the system, whose deadlines are met. The rest of the paper is structured as follows. Section 2 discusses the system model. In Section 3, related work and motivations for our work are presented. In Section 4, we propose an algorithm for fault-tolerant scheduling of real-time tasks, and also propose some enhancements to it. In Section 5, the performance of the proposed algorithm together with its enhancements is studied through simulation, and also compared with an algorithm proposed recently in [3]. Finally, in Section 6, we make some concluding remarks. System Model In this section, we first present the task model, followed by scheduler model, and then some definitions which are necessary to explain the scheduling algorithm. 2.1 Task Model 1. Tasks are aperiodic, i.e., the task arrivals are not known a priori. Every task T i has the attributes: arrival time (a i ), ready time (r i ), worst case computation time 2. The actual computation time of a task T i , denoted as c i , may be less than its worst case computation time due to the presence of data dependent loops and conditional statements in the task code, and due to architectural features of the system such as cache hits and dynamic branch prediction. The worst case execution time of a task is obtained based on both static code analysis and the average of execution times under possible worst cases. There might be cases in which the actual computation time of a task may be more than its worst case computation time. There are techniques to handle such situations. One such technique is "Task Pair" scheme [28] in which the worst case computation time of a task is added with the worst case computation time of an exception task. If the actual computation time exceeds the (original) worst case computation time, the exception task is invoked. 3. Resource constraints: A task might need some resources such as data structures, variables, and communication buffers for its execution. Each resource may have multiple instances. Every task can have two types of accesses to a resource: a) exclusive access, in which case, no other task can use the resource with it or b) shared access, in which case, it can share the resource with another task (the other task also should be willing to share the resource). Resource conflict exists between two tasks T i and T j if both of them require the same resource and one of the accesses is exclusive. 4. Each task T i has two versions, namely, primary copy and backup copy. The worst case computation time of a primary copy may be more than that of its backup. The other attributes and resource requirements of both the copies are identical. 5. Each task encounters at most one failure either due to processor failure or due to software failure, i.e., if the primary fails, its backup always succeeds. 6. Tasks are non-preemptable, i.e., when a task starts execution on a processor, it finishes to its completion. 7. Tasks are not parallelizable, which means that a task can be executed on only one processor. This necessitates the sum of worst case computation times of primary and backup copies should be less than or equal to (d that both the copies of a task can be schedulable within this interval. 8. The system has multiple identical processors which are connected through a shared medium. 9. Faults can be transient or permanent, and are independent, i.e., correlated failures are not considered 10. There exists a fault-detection mechanism such as acceptance tests to detect both processor failures and software failures. Most complex real-time applications have both periodic and aperiodic tasks. The dynamic planning based scheduling approach used in this paper is also applicable to such real-time applications as described below. The system resources (including processors) are partitioned into two sets, one for periodic tasks and the other for aperiodic tasks. The periodic tasks are scheduled by a static table-driven scheduling approach [22] onto the resource partition corresponding to periodic tasks and the aperiodic tasks are scheduled by a dynamic planning based scheduling approach [21, 22, 13] onto the resource partition corresponding to aperiodic tasks. Tasks may have precedence constraints. Ready times and deadlines of tasks can be modified such that they comply with the precedence constraints among them. Dealing with precedence constraints is equivalent to working with the modified ready times and deadlines [11]. Therefore, the proposed algorithm can also be applied to tasks having precedence constraints among them. 2.2 Scheduler Model In a dynamic multiprocessor scheduling, all the tasks arrive at a central processor called the scheduler, from where they are distributed to other processors in the system for execution. The communication between the scheduler and the processors is through dispatch queues. Each processor has its own dispatch queue. This organization, shown in Fig.1, ensures that the processors will always find some tasks (if there are enough tasks in the system) in the dispatch queues when they finish the execution of their current tasks. The scheduler will be running in parallel with the processors, scheduling the newly arriving tasks, and periodically updating the dispatch queues. The scheduler has to ensure that the dispatch queues are always filled to their minimum capacity (if there are tasks left with it) for this parallel operation. This minimum capacity depends on the worst case time required by the scheduler to reschedule its tasks upon the arrival of a new task [24]. The scheduler arrives at a feasible schedule based on the worst case computation times of tasks satisfying their timing, resource, and fault-tolerant constraints. The use of one scheduler for the whole system makes the scheduler a single point of failure. The scheduler can be made fault-tolerant by employing modular redundancy technique in which a backup scheduler runs in parallel with the primary scheduler and both the schedulers perform an acceptance test. The dispatch queues will be updated by one of the schedulers which passes the acceptance test. A simple acceptance test for this is to check whether each task in the schedule finishes before its deadline satisfying its requirements. tasks Task queue Current schedule dispatch queues Dispatch queues (Feasible schedule) Processors Scheduler Min. length of P2Fig.1 Parallel execution of scheduler and processors 2.2.1 Resource Reclaiming Resource reclaiming [24] refers to the problem of utilizing resources (processors and other resources) left unused by a task (version) when: (i) it executes less than its worst case computation time, or (ii) it is deleted from the current schedule. Deletion of a task version takes place when extra versions are initially scheduled to account for fault tolerance, i.e., in the PB fault-tolerant approach, when the primary version of a task completes its execution successfully, there is no need for the temporally redundant backup version to be executed and hence it can be deleted. Each processor invokes a resource reclaiming algorithm at the completion of its currently executing task. If resource reclaiming is not used, processors execute tasks strictly based on the scheduled start times as per the feasible schedule, which results in making the resources remain unused, thus reducing the guarantee ratio. The scheduler is informed with the time reclaimed by the reclaiming algorithm so that the scheduler can schedule the newly arriving tasks correctly and effectively. A protocol for achieving this is suggested in [24]. Therefore, any dynamic scheduling scheme should have a scheduler with associated resource reclaiming. 3 Background In this section, we first discuss the existing work on fault-tolerant scheduling, and then highlight the limitations of these works which form the motivation for our work. 3.1 Related Work Many practical instances of scheduling problems have been found to be NP-complete [2], i.e., it is believed that there is no optimal polynomial-time algorithm for them. It was shown in [1] that there does not exist an algorithm for optimally scheduling dynamically arriving tasks with or without mutual exclusion constraints on a multiprocessor system. These negative results motivated the need for heuristic approaches for solving the scheduling problem. Recently, many heuristic scheduling algorithms [21, 13] have been proposed to dynamically schedule a set of tasks whose computation times, deadlines, and resource requirements are known only on arrival. For multiprocessor systems with resource constrained tasks, a heuristic search algorithm, called myopic scheduling algorithm, was proposed in [21]. The authors of [21] have shown that the integrated heuristic used there which is a function of deadline and earliest start time of a task performs better than simple heuristics such as earliest deadline first, least laxity first, and minimum processing time first. In [10], a PB scheme has been proposed for preemptively scheduling periodic tasks in a uniprocessor system. This approach guarantees that (i) a primary copy meets its deadline if there is no failure and (ii) its backup copy will run by the deadline if there is a failure. To achieve this, it precomputes tree of schedules (where the tree can be encoded within a table-driven scheduler) by considering all possible failure scenarios of tasks. This scheme is applicable to simple periodic tasks, where the periods of the tasks are multiples of the smallest period. The objective of this approach is to increase the number of primary task executions. Another PB scheme is proposed in [19] for scheduling periodic tasks in a multiprocessor system. In this strategy, a backup schedule is created for each set of tasks in the primary schedule. The tasks are then rotated such that primary and backup schedules are on different processors and they do not overlap. This approach tolerates up to one failure in the worst case, by using double the number of processors used in the corresponding non-fault-tolerant schedule. In [7], processor failures are handled by maintaining contingency or backup schedules. These schedules are used in the event of a failure. The backup schedules are generated assuming that an optimal schedule exists and the schedule is enhanced with the addition of "ghost" tasks, which function primarily as standby tasks. The addition of tasks may not be possible in some schedules. A PB based algorithm with backup overloading and backup deallocation has been proposed recently [3] for fault-tolerant dynamic scheduling of real-time tasks in multiprocessor systems, which we call as backup overloading algorithm. The backup overloading algorithm allocates more than a single backup in a time interval (where time interval of a task is the interval between scheduled start time and scheduled finish time of the task) and deallocates the resources unused by the backup copies in case of fault-free operation. Two or more backups can overlap in the schedule (overloading) of a processor, if the primaries of these backups are scheduled on different processors. The concept of backup overloading is valid under the assumption that there can be at most one fault at any instant of time in the entire system. In [3], it was shown that backup deallocation is more effective than the backup overloading. The paper also provides a mechanism to determine the number of processors required to provide fault-tolerance in a dynamic real-time system. Discussion about other related work on fault-tolerant real-time scheduling can be found in [3]. 3.2 Motivations for Our Work The algorithms discussed in [7, 19] are static algorithms and cannot be applied to dynamic scheduling, considered in this paper, due to their high complexities. The algorithm discussed in [10] is for scheduling periodic tasks in uniprocessor systems and cannot be extended to the dynamic scheduling as it expects the tasks to be periodic. Though the algorithm proposed in [3] is for dynamic scheduling, it does not consider resource constraints among tasks which is a practical requirement in any complex real-time system, and assumes at most one failure at any instant of time, which is pessimistic. The algorithm in [3] is able to deallocate a backup when its primary is successful and uses this reclaimed (processor) time to schedule other tasks in a greedy manner. The resource reclaiming in such systems is simple and is said to be work-conserving which means that it never leaves a processor idle if there is a dispatchable task. But, resource reclaiming on multiprocessor systems with resource constrained tasks is more complicated. This is due to the potential parallelism provided by a multiprocessor and potential resource conflicts among tasks. When the actual computation time of a task differs from its worst case computation time in a non-preemptive multiprocessor schedule with resource constraints, run-time anomalies may occur [4] if a work-conserving reclaiming scheme is used. These anomalies may cause some of the already guaranteed tasks to miss their deadlines. In particular, one cannot use a work-conserving scheme for resource reclaiming from resource constrained tasks. Moreover, the algorithm proposed in [3] does not reclaim resources when the actual computation times of tasks are less than their worst case computation times, which is true for many tasks. But, resource reclaiming in such cases is very effective in improving the guarantee ratio [24]. The Spring scheduling approach [27] schedules dynamically arriving tasks with resource requirements and reclaims resources from early completing tasks and does not address the fault-tolerant requirements explicitly. Our algorithm works within the Spring scheduling approach and builds fault-tolerant solutions around it to support PB based fault-tolerance. To the best of our knowledge, ours is the first work which addresses the fault-tolerant scheduling problem in a more practical model, which means that our algorithm handles resource constraints among tasks and reclaims resources both from early completing tasks and deallocated backups. The performance of our algorithm is compared with the backup overloading algorithm in Section 5.5. 4 The Fault-tolerant Scheduling Algorithm In this section, we first define some terms and then present our fault-tolerant scheduling algorithm which uses these terms. 4.1 Terminology 1: The scheduler fixes a feasible schedule S. The feasible schedule uses the worst case computation time of a task for scheduling it and ensures that the timing, resource, and fault-tolerant constraints of all the tasks in S are met. A partial schedule is one which does not contain all the tasks. Definition 2: st(T i ) is the scheduled start time of task T i which satisfies r the scheduled finish time of task T i which satisfies r Definition 3: EAT s k ) is the earliest time at which the resource R k becomes available for shared (exclusive) usage [21]. Definition 4: Let P be the set of processors, and R i be the set of resources requested by task T i . Earliest start time of a task T i , denoted as EST(T i ), is the earliest time when its execution can be started, which is defined as: EST (T i denotes the time at which the processor P j is available for executing a task, and the third term denotes maximum among the available time of the resources requested by task T i , in which shared mode and exclusive mode. Definition 5: proc(T i ) is the processor to which task T i is scheduled. The processor to which task T i should not get scheduled is denoted as exclude proc(T i ). Definition is the scheduled start time and f t(Pr i ) is the scheduled finish time of primary copy of a task T i . Similarly, st(Bk i ) and f t(Bk i ) denote the same for the backup copy of T i . Definition 7: The primary and backup copies of a task T i are said to be mutually exclusive in time, denoted as time exclusion(T i Definition 8: The primary and backup copies of a task T i are said to be mutually exclusive in space, denoted as space exclusion(T i A task is said to be feasible in a fault-tolerant schedule if it satisfies the following conditions: ffl The primary and backup copies of a task should satisfy: r i - st(Pr This is because both the copies of a task must satisfy the timing constraints and it is assumed that the backup is executed after the failure in its primary is detected (time exclusion). Failure detection is done through acceptance test or some other means only at the completion of every primary copy. The time exclusion between primary and backup copies of a task can be relaxed if the backup is allowed to execute in parallel [5, 30] (or overlap) with its primary. This is not preferable in dynamic scheduling as discussed below. ffl Primary and backup copies of a task should mutually exclude in space in the schedule. This is necessary to tolerate permanent processor failures. Mutual exclusion in time is very useful from the resource reclaiming point of view. If the primary is successful, the backup need not be executed and the time interval allocated to the backup can be reclaimed fully, if primary and backup satisfy time exclusion, thereby improving the schedulability [15]. In other words, if primary and backup of a task overlap in execution, the backup unnecessarily executes (in part or full) even when its primary is successful. This would result in poor resource utilization, thereby reducing the schedulability. Moreover, overlapping of primary and backup of a task introduces resource conflicts (if the access is exclusive) among them since they have the same resource requirements and forces them to exclude in time if only one instance of the requested resource is available at that time. 4.2 The Distance Myopic Algorithm The Spring scheduling [27] approach uses a heuristic search algorithm (called myopic algorithm [21]) for non-fault-tolerant scheduling of resource constrained real-time tasks in a multiprocessor system, and uses Basic or Early start algorithms for resource reclaiming. One of the objectives of our work here is to propose fault-tolerant enhancements to the Spring scheduling approach. We make the following enhancements to the Spring scheduling to support PB based fault-tolerance: ffl a notion of distance is introduced, which decides the relative difference in position between primary and backup copies of a task in the task queue. ffl flexible level of backup overloading; this introduces a tradeoff between number of faults in the system and the system performance. ffl use of restriction vector (RV) [15] based algorithm to reclaim resources from both deallocated backups and early completing tasks. 4.2.1 Notion of Distance Since in our task model, every task, T i , has two copies, we place both of them in the task queue with relative difference of Distance(Pr positions. The primary copy of any task always precedes its backup copy in the task queue. Let n be the number of currently active tasks whose characteristics are known. The algorithm does not know the characteristics of new tasks, which may arrive, while scheduling the currently active tasks. The distance is an input parameter to the scheduling algorithm which determines the relative positions of the copies of a task in the task queue in the following way: distance for the first (n \Gamma (n mod distance)) tasks mod distance for the last (n mod distance) tasks, The following is an example task queue with assuming that the deadlines of tasks are in the non-decreasing order. The positioning of backup copies in the task queue relative to their primaries can easily be achieved with minimal cost: (i) by having two queues, one for primary copies (n entries) and the other for backup copies (n entries), and (ii) merging these queues, before invoking the scheduler, based on the distance value to get a task queue of 2n entries. The cost involved due to merging is 2n. 4.2.2 Myopic Scheduling Algorithm The myopic algorithm [21] is a heuristic search algorithm that schedules dynamically arriving real-time tasks with resource constraints. It works as follows for scheduling a set of tasks. A vertex in the search tree represents a partial schedule. The schedule from a vertex is extended only if the vertex is strongly feasible. A vertex is strongly feasible if a feasible schedule can be generated by extending the current partial schedule with each task of the feasibility check window. Feasibility check window is a subset of first K unscheduled tasks. Larger the size of the feasibility check window, higher the scheduling cost and more the look ahead nature. If the current vertex is strongly feasible, the algorithm computes a heuristic function, for each task within the feasibility check window, based on deadline and earliest start time of the task. It then extends the schedule by the task having the best (smallest) heuristic value. Otherwise, it backtracks to the previous vertex and then the schedule is extended from there using a task which has the next best heuristic value. 4.2.3 The Distance Based Fault-tolerant Myopic Algorithm We make fault-tolerant extensions to the original myopic algorithm using the distance concept for scheduling a set of tasks. Here, we assume that each task is a plan. The algorithm attempts to generate a feasible schedule for the task set with minimum number of rejections. Distance Myopic() 1. Order the tasks (primary copies) in non-decreasing order of deadlines in the task queue and insert the backup copies at appropriate distance from their primary copies. 2. Compute Earliest Start Time EST (T i ) for the first K tasks, where K is the size of the feasibility check window. 3. Check for strong feasibility: check whether EST (T i is true for all the K tasks. 4. If strongly feasible or no more backtracking is possible (a) Compute the heuristic function for the first K tasks, where W is an input parameter. ffl When Bk i of task T i is considered for H function evaluation, if Pr i is not yet scheduled, set EST (Bk i (b) Choose the task with the best (smallest) H value to extend the schedule. (c) If the best task meets its deadline, extend the schedule by the best task (best task is accepted in the schedule). ffl If the best task is primary copy (Pr i ) of task T i This is to achieve time exclusion for task T i . This is to achieve space exclusion for task T i . (d) else reject the best task and move the feasibility check window by one task to the right. (e) If the rejected task is a backup copy, delete its primary copy from the schedule. 5. else Backtrack to the previous search level and try extending the schedule with a task having the next best H value. 6. Repeat steps (2-5) until termination condition is met. The termination condition is either (i) all the tasks are scheduled or (ii) all the tasks are considered for scheduling and no more backtrack is possible. The complexity of the algorithm is the same the original myopic algorithm, which is O(Kn). It is to be noted that the distance myopic algorithm can tolerate more than one fault at any point of time, and the number of faults is limited by the assumption that at most one of the copies of a task can fail. Once a processor fault is detected, the recovery is inherent in the schedule meaning that the backups, of the primaries scheduled on the failed processors, will always succeed. In addition, whether the failed processors will be considered or not for further scheduling depends on the type of fault. If it is a transient processor fault, the processor on which the failure has occurred will be considered for further scheduling. On the other hand, if it is a permanent processor fault, the processor on which the failure has occurred will not be considered for further scheduling till it gets repaired. If the failure is due to task error (software fault), it is treated like a transient processor fault. 4.2.4 Flexible Backup Overloading in Distance Myopic Here, we discuss as how to incorporate flexible level of backup overloading into the distance myopic algorithm. This introduces a tradeoff between the number of faults in the system and the guarantee ratio. Before, defining the flexible backup overloading, we state from [3] the condition under which backups can be overloaded. If Pr i and Pr j are scheduled on two different processors, then their backups Bk i and Bk j can overlap in execution on a processor: The backup overloading is depicted in Fig.2. In Fig.2, Bk 1 and Bk 3 which are scheduled on processor in execution, whose primaries Pr 1 and Pr 3 are scheduled on different processors P 1 and P 3 , respectively. This backup overloading is valid under the assumption that there is at most one failure, in the system (at any instant of time). This is a too pessimistic assumption especially when the number of processors in the system is large. Processor 1 Processor 2 Processor 3 Primary 1 Primary 2 Primary 4 Primary 3 Time Backups 1 and 3 Fig.2 Backup overloading We introduce flexibility in overloading (and hence the number of faults) by forming the processors into different groups. Let group(P i ) denote the group in which processor P i is a member, and m be the number of processors in the system. The rules for flexible backup overloading are: Every processor is a member of exactly one group. ffl Each group should have at least three processors for backup overloading to take place in that group. ffl Size of each group (gsize) is the same, except for one group, when (m=gsize) is not an integer. ffl Backup overloading can take place only among the processors in a group: ffl Both primary and backup copies of a task are to be scheduled on to the processors of the same group. The flexible overloading scheme permits at most d(m=gsize)e number of faults at any instant of time, with a restriction that there is at most one fault in each group. In the flexible overloading scheme, when the number of faults permitted is increased, the flexibility in backup overloading is limited and hence the guarantee ratio might drop down. This mechanism gives the flexibility for the system designer to choose the desired degree of fault-tolerance. In Section 5.2.5, we study the tradeoff between the number of faults and the performance of the system. 4.2.5 Restriction Vector Based Resource Reclaiming In our dynamic fault-tolerant scheduling approach, we have used restriction vector (RV) algorithm for resource reclaiming. RV algorithm uses a data structure called restriction vector which captures resource, precedence, and fault-tolerant constraints among tasks in a unified way. Each task T i has an associated m-component vector, RV i [1::m], called Restriction Vector, where m is the number of processors in the system. RV i [j] for a task T i contains the last task in T !i (j), where T !i (j) is the set of tasks which are scheduled to finish before T i begins. Before updating the dispatch queues, the scheduler computes the restriction vector for each of the tasks in the feasible schedule. For computing RV of a task T i , the scheduler checks at most k tasks (in the order of latest finish time first) which are scheduled to finish on other processors before T i starts execution. The latest task on processor j which has resource conflict or precedence relation with the task T i becomes RV i [j]. If no such task exists, then the k-th task is RV i [j]. The RV algorithm [15] says: start executing a task T i only if processor on which T i is scheduled is idle and all the tasks in its restriction vector have successfully finished their execution. Performance Studies In this section, we first present the simulation studies on various algorithms, and then present an analytical study based on Markov chains which highlights the importance of distance parameter in fault-tolerant dynamic scheduling. The simulation experiments were conducted in two parts. The first part highlights the importance of distance parameter and the second part highlights the importance of each of the guarantee ratio improving techniques, namely, distance concept, backup deallocation, and backup overloading. For each point in the performance curves (Figs.4-15), the total number of tasks arrived in the system is 20,000. The parameters used in the simulation studies are given in Fig.3. The tasks for the simulation are generated as follows: 1. The worst case computation times of primary copies are chosen uniformly between MIN C and MAX C. 2. The deadline of a task T i (primary copy) is uniformly chosen between r 2. 3. The inter arrival time of tasks (primary copies) is exponentially distributed with mean 1=(- m) (MIN C +MAX C)=2. 4. The actual computation time of a primary copy is chosen uniformly between MIN C and its worst case computation time, if aw-ratio is random (rand). Otherwise, it is aw-ratio times the worst case computation time. 5. The backup copies are assumed to have identical characteristics of their primary copies. parameter explanation value taken when (varied) (fixed) MIN C minimum computation time of tasks (-) (40) MAX C maximum computation time of tasks (-) (80) - task arrival rate (0.2,0.3,.,0.7) (0.5 or 0.4) R laxity parameter (2,3,.,7) (4) UseP probability that a task uses (0.1,0.2,.,0.5) (0.4) a resource ShareP probability that a task accesses (-) (0.4) a resource in shared mode K size of feasibility check window (1,3, ., 11) (3) W weight of EST(T i ) in the H function (-) (1) aw-ratio ratio of actual to worst case (0.5,0.6,.,1.0) (rand) computation times FaultP probability that a primary fails (0.1,0.2,.,0.5) (0.2) distance relative difference in positions of primary (1,5,9,13) (5) and backup copies in the task queue m number of processors (5,6,.,10) (8) NumRes number of resources (-) (2) NumInst number of instances per resource (-) (2) Fig.3 Simulation parameters 5.1 Experiments Highlighting Distance Parameter In this section, we present the simulation results obtained for different values of distance parameter by varying K, -, UseP , and FaultP parameters. For this study, the - value is taken as 0.5 when fixed. The algorithms studied here reclaims resources both from early completing tasks and deallocated backups. The actual computation time of a task is chosen uniformly between MIN C and its worst case computation time. 5.1.1 Effect of Feasibility Check Window Fig.4 shows the effect of varying distance for different values of K. Note that for larger values of K, the number of H function evaluations and EST() computations are also more, which means higher scheduling cost. The interplay between the distance and size of the feasibility check window is described below. ffl When distance is small, the position of backup copies in the task queue is close to their respective primary copies and hence the possibility of scheduling these backup copies may get postponed (we call this, backup postponement) due to time and space exclusions. This makes more and more unscheduled backup copies getting accumulated. When this number exceeds K, the scheduler is forced to choose the best task (say T b ) among these backup copies, which results in creation of a hole 1 in the schedule since EST(T b ) is greater than the available time (avail time) of idle processors. This hole creation can be avoided by moving the feasibility check window till a primary task falls into it. However, we do not consider this approach since it increases the scheduling cost. ffl When distance is large, the position of the backup copies in the task queue is far apart from their respective primary copies, i.e., tasks (backup copies) having lower deadlines may be placed after some tasks (primary copies) having higher deadlines. This may lead to backtracks (and hence rejection, if no backtrack is possible) when the feasibility check window reaches these backup copies (we call this, forced backtrack). The guarantee ratio increases with increasing K for a given distance for some time (growing phase) and then starts decreasing for higher values of K (shrinking phase). From Fig.4, the shrinking phase starts at K values 7,5,5, and 7, for distance values 1,5,9, and 13, respectively. The reason for this is that the backup postponement is very high at the beginning of the growing phase, decreases along with it and reaches the lowest value at the end of it (equivalently, beginning of the shrinking phase), and the number of forced rejections is very low at the beginning of the shrinking phase and increases along with it. This reveals two facts: (a) increased value of K (increased look ahead nature) does not necessarily increase the guarantee ratio and (b) the optimal K for each distance is different. The right combination of K and distance offers the best guarantee ratio. From Fig.4, the best guarantee ratio is obtained when 9. Suppose two distance values give the same (best) guarantee ratio, the one with lower K is preferable because of lower scheduling cost. 5.1.2 Effect of Resource Usage, Task Load, and Fault Probability In Fig.5, the probability that a task uses a resource (UseP ) is varied. For a fixed value of ShareP (= 0.4), higher UseP means more resource conflicts among tasks. From Fig.5, it can be seen that the guarantee ratio decreases as UseP increases. This is applicable for all values of distance. From Fig.5, for most of the values of UseP , better guarantee ratio is obtained when distance is 9. The task arrival rate has been varied in Fig.6. Higher - means lower inter arrival time and hence higher task load. From Fig.6, it can be seen that increasing - decreases the guarantee ratio for all values unusable processor time for scheduling. of distance. From Fig.6, for most of the values of -, better guarantee ratio is obtained when distance is 5 and 9 compared to other values. In Fig.7, the probability that a primary copy encounters a failure is varied. As FaultP increases, the guarantee ratio decreases. This is applicable for all values of distance. From Fig.7, when distance is 5 and 9, better guarantee ratio is obtained compared to the other values of distance.50602 4 Guarantee ratio Size of feasibility check window Fig.4 Effect of feasibility check window5060708090 Guarantee ratio Resource usage probability Fig.5 Effect of resource usage probability507090 Guarantee ratio Task arrival rate Fig.6 Effect of task load52566064 Guarantee ratio Primary fault probability Fig.7 Effect of primary fault probability 5.1.3 Choice of Distance Based on the observations of simulation studies, a simple heuristic to select good K and distance value is based on the number of processors, and the number of resources and their usage. If there are few resources with high UseP and low ShareP , then there exists more resource conflicts among tasks. In such cases, the EST() of a task is mostly decided by the resource available time rather than processor available time or task ready time. Therefore, large value of K might help in such situations. The value of distance may be approximately equal to a value in the range [m/2, m] since at most m consecutive primaries or backups can be scheduled in the worst case. The value of K may be less than the distance since larger K means higher scheduling cost, which might nullify or reduce the gain obtained. 5.2 Experiments Highlighting GR Improving Techniques In this section, we show through simulation the importance of each of the guarantee ratio (GR) improving techniques, namely, distance concept, backup deallocation, and backup overloading. For this experiments, we have taken the - value to be 0.4, when fixed. The actual computation time of a task is chosen uniformly between MIN C and its worst case computation time. The plots in Figs.8-13 correspond to the following algorithms: Myopic. This is a fault-tolerant version of myopic algorithm with distance = 1. This algorithm reclaims resources only from early completing tasks. Distance concept. This is same as algorithm A0 except that distance = 5 (this value for distance is chosen based on the previous experiments and discussions). deallocation. This is algorithm A1 together with resource reclaiming from deallocated backups as well. overloading. This is algorithm A2 together with backup overloading. For this full overloading is considered, i.e., gsize = m. This algorithm permits at most one failure, whereas the other algorithms can tolerate more than one failure. The difference in guarantee ratio between algorithms: (i) A0 and A1 is due to distance concept, (ii) A1 and A2 is due to backup deallocation, and (iii) A2 and A3 is due to backup overloading. From Figs.8-13, it can be see that each of the guarantee ratio improving techniques improves the guarantee ratio of the system, with very minimal increase in scheduling cost. That is, algorithms A0, A1, A2, and A3 offer non-decreasing order of guarantee ratio. The distance concept and backup deallocation are more effective compared to backup overloading. 5.2.1 Effect of Task Laxity, Resource Usage, and Task Load The effect of task laxity (R) is studied in Fig.8. As the laxity increases, the guarantee ratio also increases. For lower laxities, the difference in guarantee ratio between algorithms is less and increases with increasing laxity. This is because, for lower values of laxity, the deadlines of tasks are very tight and due to which the guarantee improving techniques have less flexibility to be more effective. In Fig.9, the probability that a task uses a resource (UseP ) is varied. The increase in UseP , for a fixed ShareP , increases the resource conflicts among tasks and hence the guarantee ratio decreases. This is true for all the algorithms. The effect of task load is studied in Fig.10. As load increases, the guarantee ratio decreases for all the algorithms. For lower values of task loads (when to 0:3), the guarantee ratio of all the four algorithms is more or less the same. This is because, at such low load, the system has enough processors and resources to feasibly schedule the tasks. When the load increases, the difference in guarantee between algorithms also increases, which means that the proposed techniques are effective at higher loads. 5.2.2 Effect of Number of Processors The effect of varying the number of processors (m) is studied in Fig.11. For this, the task load is fixed to be the load of 8 processors. The increase in number of processors increases the guarantee ratio for all the four algorithms. The difference in guarantee ratio for two successive values of m (i.e., m and m+ 1) is very high when m is small, and decreases as m increases. This is because of limited availability of resources, i.e., the bottleneck is the resources and not the processors. This means that if m is increased beyond 10, there cannot be much improvement in the guarantee ratio.6670747882 Guarantee ratio Task laxity A3 Fig.8 Effect of task laxity65758595 Guarantee ratio Resource usage probability A3 Fig.9 Effect of resource usage probability Guarantee ratio Task arrival rate A3 Fig.10 Effect of task load50607080 Guarantee ratio Number of processors A3 Fig.11 Effect of number of processors7072747678 Guarantee ratio Actual to worst case computation ratio A3 Fig.12 Effect of actual to worst case computation Guarantee ratio Primary fault probability Fig.13 Effect of primary fault probability 5.2.3 Effect of Actual to Worst Case Computation Time Ratio The ratio between actual to worst case computation time (aw-ratio) of tasks is varied in Fig.12. In this experiment, the actual computation time of a task is taken to be aw-ratio times the worst case computation of the task. From Fig.12, an increase in aw-ratio decreases the guarantee ratio for all the algorithms. When aw-ratio=1.0, the reclaiming is only due to backup deallocation (wherever applicable). For example, for algorithms A0 and A1, when aw-ratio=1.0, no resource reclaiming is possible. When aw-ratio=1.0, the difference in guarantee ratio between A0 and A1 is purely due to distance concept, between A1 and A2 is purely due to backup deallocation, and between A2 and A3 is purely due to backup overloading. 5.2.4 Effect of Fault Probability The probability that a primary encounters a fault (FaultP ) is varied in Fig.13. Here, only three algorithms (A0, A1, and A2) are plotted because the number of faults (for a given FaultP ) generated while studying A3 is different (very less), because of at most one fault at a time, compared to the other algorithms. When there is no fault in the system, which means that every backup is deallocated. The guarantee ratio of algorithms A0 and A1 is flat for all values of FaultP since they do not deallocate backups. For A2, an increase in FaultP decreases the guarantee ratio. 5.2.5 Performance of Flexible Overloading The performance of flexible backup overloading has been studied for various parameters. Here, we present only some sample results. For these experiments, m is taken as 8, and the different algorithms studied are: (i) no overloading (Algorithm A2), (ii) half overloading (gsize = (say A4), and (iii) full overloading which is the same as algorithm A3. The tradeoff between performance and fault-tolerance is studied through this experiment. At any point of time, Algorithm A2 can tolerate more than one fault, algorithm A4 can tolerate two faults with a restriction that there is at most one fault within a group, and algorithm A3 can tolerate at most one fault. The task load and laxity are varied in Fig.14 and 15, respectively. From these figures, the guarantee ratio offered by full overloading is better than the other two, and half overloading is better than no over- loading. The gain in guarantee ratio obtained by trading (reducing) the number faults in full overloading is around 2% to 3% in both the experiments. For lower task loads, the gain is less than 1% and is more at higher task loads. This reveals that backup overloading is less effective in improving the guarantee ratio compared to the other techniques such as distance concept, backup deallocation, and reclaiming from early completing tasks. Thus, the flexible overloading provides a tradeoff between performance and the degree of fault-tolerance. Guarantee ratio Task arrival rate A3 Fig.14 Effect of task load70747882 Guarantee ratio Task laxity A3 Fig.15 Effect of task laxity 5.3 Analytical Study In this section, we show analytically using Markov chains that the distance is an important parameter in fault-tolerant dynamic scheduling of real-time tasks. Using Markov chains, the possible states of the system and the probabilities of transitions among them can be determined, which can then be used to evaluate different dependability metrics for the system. The analysis presented here is similar to the one given in [3, 17] except for the backup preallocation strategy. To make the analysis tractable, we make the following assumptions: 1. All tasks have unit worst case computation time, i.e., c 2. Backup slots are preallocated in the schedule based on the distance parameter. 3. FIFO scheduling strategy is used. 4. Size of the feasibility check window (K) is 1. 5. Task deadlines are uniformly distributed in the interval [W min ,W max ] relative to their ready times, which we call deadline window. 6. Task arrivals are uniformly distributed with mean A av . 7. Backup overloading and resource reclaiming are not considered. 5.3.1 Backup Preallocation Strategy Since the tasks are of unit length, we reserve slots in the schedule for the backup copies based on the distance parameter. Let m be the number of processors and d be the distance. Let s and In our backup preallocation strategy, at any time t, the available number of primary slots is s 1 if t is odd, s 2 if t is even. Similarly, the available number of backup slots is s 2 if t is odd, s 1 if t is even. In other words, backups are reserved at time slot the primaries of time slot t. Fig.16 shows the backup preallocation with 2. Note that the backup preallocation for m processors with distance (m \Gamma d) is the same as for m processors with distance d. In our backup preallocation strategy, d should not be equal to (m \Gamma 1) because if a primary is scheduled on slot t (even), its backup slot is already reserved on the same processor at time slot which is a violation of space exclusion. Also, d ? m does not have any meaning in the preallocation. Processors Time (t) -> Backup slot Fig.16 Distance based backup preallocation with 5.4 Analysis If P ar (k) is the probability of k tasks arriving at a given time, then P ar If Pwin (k) is the probability that an arriving task has a relative deadline w, then The arriving tasks (primary copies) are appended to the task queue (Q) and they are scheduled in FIFO order. Given that s 1 or s 2 tasks can be scheduled on a given time slot t depending on whether t is odd or even, respectively, then the position of the tasks in the Q indicates their scheduled start times. If at the beginning of time slot t, a task T i is the k-th task in Q, then T i is scheduled to execute at time k is the time, from now, at which a task will execute whose position in the Q is k and is defined as In equation (3), arrives at time t, its schedulability depends on the length of Q and on the relative deadline w i of the task. If T i is appended at position q of Q and w i - g q , then the primary copy, Pr i , is guaranteed to execute before time the task is not schedulable since it will miss its deadline. Moreover, if w guaranteed to execute before t +w i . Note that in our backup preallocation strategy, the backup of a task is scheduled in the immediate next slot of its primary. The dynamics of the system can be modelled using Markov chain in which each state represents the number of tasks in Q and each transition represents the change in the length of the Q in one unit of time. The probabilities of different transitions may be calculated from the rate of task arrival. For simplicity, the average number of tasks executed at any time t is which is m=2. If S u represents the state in which Q contains u tasks and u - m=2, then the probability of a transition from S u to S u\Gammam=2+k is P ar (k) since at any time t, k tasks can arrive and m=2 tasks get executed. If only u tasks are executed, then there is a state transition from S u to S k with probability ar (k). When the k arriving tasks have finite deadlines, some of these tasks may be rejected. Let P q;k be the probability that one of the k tasks is rejected when the queue size is q. The value of P q;k is the probability that the relative deadline of the task is smaller than and the extra one time unit is needed to schedule the backup. Then, Pwin (w); (4) Hence, when the queue size is q, the probability, P rej (r; k; q), that r out of the k tasks are rejected is r is the number of possible ways to select r out of k elements. Our objective is to find the guarantee ratio (rejection ratio) for different values of distance. To do that, we need to compute the number of tasks rejected in each state. This is done by splitting each state S u in the one-dimensional Markov chain into 2A av av is the maximum number of task arrivals, and possibly rejected, in unit time. In the two-dimensional Markov chain, the state S u;r represents the queue size as u and r tasks were rejected when the transition was made into S u;r . The two-dimensional Markov chain contains (m=2)W columns (number of arrivals in unit time 1), and the transition probabilities become: if if By computing the steady state probabilities of being in the rejection states, it is possible to compute the expected value of the number of rejected tasks Rej per unit time. If P ss (u; v) is the steady state probability of being in state S u;v , then (vP ss (u; v)): (6) Then, the rate of task rejection is given by Rej=A av . Note that, P ss (u; 0) is not included in equation (6) since these are the states corresponding to no rejection. 5.4.1 Results Figs.17 and show the rejection ratio by varying distance for different values of A av and W max , re- spectively. The values of the other fixed parameters are also given in the figures. Since the preallocation of backups for distance d and (m \Gamma d) is identical, their corresponding rejection ratios are also the same. From the plots, it can be observed that the rejection ratio varies with varying distance. For lower values of distance, the rejection ratio is more and the same is true for higher values of distance. The lowest rejection ratio (best guarantee ratio) corresponds to some medium value of distance. From the Figs.17- 19, the optimal value of distance is m=2. Therefore, the distance parameter plays a crucial role on the effectiveness of dynamic fault-tolerant scheduling algorithms. Rejection ratio Distance Fig.17 Effect of task load Rejection ratio Distance Fig.18 Effect of laxity Rejection ratio Distance Fig.19 Effect of distance 5.5 Comparison with an Existing Algorithm In this section, we compare our distance myopic algorithm with a recently proposed algorithm by Ghosh, Melhem, and Moss~e (which we call, GMM algorithm) in [3] for fault-tolerant scheduling of dynamic real-time tasks. The GMM algorithm uses full backup overloading (gsize = m) and backup deallocation, and permits at most one failure at any point of time. The GMM algorithm does not address resource constraints among tasks and reclaims resources only due to backup deallocation. The limitations of this algorithm have been discussed in Section 3.2. In the GMM algorithm, the primary and backup copies of a task are scheduled in succession. In other words, the distance is always 1. The algorithm is informally stated below: GMM Algorithm(): begin 1. Order the tasks in non-decreasing order of deadline in the task queue. 2. Choose the first (primary) and second (backup) tasks for scheduling: ffl Schedule the primary copy as early as possible by End Fitting() or Middle Fitting() or Middle Adjusting(). ffl Schedule the backup copy as late as possible by Backup Overloading() or End Fitting() or Middle Fitting() or Middle Adjusting(). 3. If both primary and backup copies meet their deadline, accept them in the schedule. 4. else reject them. (a) End Fitting(): Schedule the current task as the last task in the schedule of a processor. (b) Middle Fitting(): Schedule the current task some where in the middle of the schedule of a processor. (c) (b) Middle Adjusting(): Schedule the current task some where in the middle of the schedule of a processor by changing start and finish times of adjacent tasks. (d) Backup Overloading(): Schedule the current task on a backup time interval if the primary copies corresponding to these backup copies are scheduled on two different processors. Each of steps (b)-(d), the search for fitting, adjusting, and overlapping, begins at the end of the schedule and proceeds towards the start of the schedule of every processor. The depth of the search is limited to an input parameter K. Since each of steps (b)-(d) takes time Km, the worst case time taken to schedule a primary copy is 2Km, whereas it is 3Km for a backup copy. The performance of distance myopic algorithms is compared with the GMM algorithm. For the sake of comparison with the GMM algorithm, no resource constraints among tasks are considered. To make the comparison fair, resource reclaiming only due to backup deallocation is considered, since GMM does not reclaim resources from early completing tasks. The plots in Figs.20 and 21 correspond to four algorithms: (i) distance myopic (DM), (ii) distance myopic with full backup overloading (DM algorithm without backup overloading (GMM - overload), and (iv) GMM algorithm. The scheduling cost of both the algorithms is made equal by appropriately setting K(= 4) and K(= 1) parameters in distance myopic and GMM algorithms, respectively. For these experiments, the values of R, UseP , FaultP , aw-ratio, and distance values are taken as 5, 0, 0.2, 1, and 5, respectively. We present here only the sample results. The task load is varied in Fig.20. In this figure, the different algorithms are ordered in decreasing order of the guarantee ratio offered: DM overloading. In Fig.21, the number of processors is varied by fixing the task load equal to the load of 8 processors. For lower number of processors, even DM algorithm is better than GMM. From these simulation experiments, we have shown that our proposed algorithm (DM + overloading) is better than the GMM algorithm even for the (restricted) task model for which it was proposed.5070900.5 0.6 0.7 Guarantee ratio Task arrival rate GMM DM Fig.20 Effect of task load305070903 4 5 6 7 8 Guarantee ratio Number of processors GMM DM Fig.21 Effect of number of processors 6 Conclusions In this paper, we have proposed an algorithm for scheduling dynamically arriving real-time tasks with resource and primary-backup based fault-tolerant requirements in a multiprocessor system. Our algorithm can tolerate more than one fault at a time, and employs techniques such as distance concept, flexible backup overloading, and resource reclaiming to improve the guarantee ratio of the system. Through simulation studies and also analytically, we have shown that the distance is a crucial parameter which decides the performance of any fault-tolerant dynamic scheduling in real-time multiprocessor systems. Our simulation studies on distance parameter show that increasing the size of feasibility check window (and hence the look ahead nature) does not necessarily increase the guarantee ratio. The right combination of K and distance offers the best guarantee ratio. We have also discussed as how to choose this combination. We have quantified the effectiveness of each of the proposed guarantee ratio improving techniques through simulation studies for a wide range of task and system parameters. Our simulation studies show that the distance concept and resource reclaiming, due to both backup deallocation and early completion of tasks, are more effective in improving the guarantee ratio compared to backup overloading. The flexible backup overloading introduces a tradeoff between the number of faults and the guarantee ratio. From the studies of flexible backup overloading, the gain (in guarantee ratio) obtained by favouring performance (i.e., reducing the number of faults) is not very significant. This indicates that the backup overloading is less effective, compared to the other techniques. We have also compared our algorithm with a recently proposed [3] fault-tolerant dynamic scheduling algorithm. Although our algorithm takes into account resource constraints among tasks and tolerates more than one fault at a time, for the sake of comparison, we restricted the studies to independent tasks with at most one failure. The simulation results show that our algorithm, when it is used with backup overloading, offers better guarantee ratio than that of the other algorithm for all task and system parameters. Currently, we are investigating as how to integrate different fault-tolerant approaches namely, triple modular redundancy, primary-backup approach, and imprecise computation into a single scheduling framework. --R "Multiprocessor on-line scheduling of hard real-time tasks," "Computers and intractability, a guide to the theory of NP- completeness," "Fault-Tolerance through scheduling of aperiodic tasks in hard real-time multiprocessor systems," "Bounds on multiprocessing timing anomalies," "Approaches to implementation of reparable distributed recovery block scheme," "Distributed fault-tolerant real-time systems," "On scheduling tasks with quick recovery from failure," "Real-time Systems," "Architectural principles for safety-critical real-time applications," "A fault tolerant scheduling problem," "Imprecise computations," "Modular redundancy in a message passing system," "An efficient dynamic scheduling algorithm for multiprocessor real-time systems," "A new study for fault-tolerant real-time dynamic scheduling algorithms," "New algorithms for resource reclaiming from precedence constrained tasks in multiprocessor real-time systems," "Real-time System Scenarios," "Analysis of a fault-tolerant multiprocessor scheduling al- gorithm," "Adaptive software fault tolerance policies with dynamic real-time guarantees," "Multiprocessor support for real-time fault-tolerant scheduling," "An environment for developing fault-tolerant software," "Efficient scheduling algorithms for real-time multiprocessor systems," "Scheduling algorithms and operating systems support for real-time systems," "Graceful degradation in real-time control applications using (m,k)-firm guarantee," "Resource reclaiming in multiprocessor real-time systems," "Real-time computing: A new discipline of computer science and engineering," "Understanding fault-tolerance and reliability," "The Spring Kernel: A new paradigm for real-time operating systems," "TaskPair-Scheduling: An approach for dynamic real-time systems," "Low overhead multiprocessor allocation strategies exploiting system spare capacity for fault detection and location," "Fault-tolerant scheduling algorithm for distributed real-time systems," "Determining redundancy levels for fault tolerant real-time systems," "Multiprocessor scheduling of processes with release times, deadlines, precedence and exclusion constraints," "Scheduling tasks with resource requirements in hard real-time systems," --TR --CTR Wei Sun , Chen Yu , Xavier Defago , Yuanyuan Zhang , Yasushi Inoguchi, Real-time Task Scheduling Using Extended Overloading Technique for Multiprocessor Systems, Proceedings of the 11th IEEE International Symposium on Distributed Simulation and Real-Time Applications, p.95-102, October 22-26, 2007 R. Al-Omari , A. K. Somani , G. Manimaran, An adaptive scheme for fault-tolerant scheduling of soft real-time tasks in multiprocessor systems, Journal of Parallel and Distributed Computing, v.65 n.5, p.595-608, May 2005 R. Al-Omari , Arun K. Somani , G. Manimaran, Efficient overloading techniques for primary-backup scheduling in real-time systems, Journal of Parallel and Distributed Computing, v.64 n.5, p.629-648, May 2004 Xiao Qin , Hong Jiang, A novel fault-tolerant scheduling algorithm for precedence constrained tasks in real-time heterogeneous systems, Parallel Computing, v.32 n.5, p.331-356, June 2006 Wenjing Rao , Alex Orailoglu , Ramesh Karri, Towards Nanoelectronics Processor Architectures, Journal of Electronic Testing: Theory and Applications, v.23 n.2-3, p.235-254, June 2007

run-time anomaly;dynamic scheduling;fault tolerance;safety critical application;resource reclaiming;real-time system

293004

New Perspectives in Turbulence.

Intermittency, a basic property of fully developed turbulent flow, decreases with growing viscosity; therefore classical relationships obtained in the limit of vanishing viscosity must be corrected when the Reynolds number is finite but large. These corrections are the main subject of the present paper. They lead to a new scaling law for wall-bounded turbulence, which is of key importance in engineering, and to a reinterpretation of the Kolmogorov--Obukhov scaling for the local structure of turbulence, which has been of paramount interest in both theory and applications. The background of these results is reviewed, in similarity methods, in the statistical theory of vortex motion, and in intermediate asymptotics, and relevant experimental data are summarized.

Introduction In February 1996 I had the privilege of meeting Prof. G.I. Barenblatt, who had just arrived in Berkeley. In our first extended conversation we discovered that we had been working on similar problems with different but complementary tools, which, when wielded in unison, led to unexpected results. We have been working together ever since, and it is a pleasure to be able to present some of the results of our joint work at this distinguished occasion. The present talk will consist of three parts: (i) An application of advanced similarity methods and vanishing-viscosity asymptotics to the analysis of wall-bounded turbulence, (ii) a discussion of the local structure of turbulence with particular attention to the higher-order structure functions, and (iii) a discussion of a near-equilibrium statistical theory of 1 Supported in part by the Applied Mathematical Sciences subprogram of the Office of Energy Research, U.S. Department of Energy, under contract DE-AC03-76-SF00098, and in part by the National Science Foundation under grants DMS94-14631 and DMS89-19074. turbulence, which motivates and complements our reading of the numerical and experimental data. The basic premise is that, as the viscosity tends to zero and the solutions of the Navier-Stokes equations acquire poorly understood temporal and spatial fluctuations, certain mean properties of the of the flow can be seen to take on well-defined limits, which can be found by expansion in a small parameter that tends to zero, albeit slowly, as the viscosity tends to zero. In the case of wall-bounded turbulence, our argument and the data show that the classical von K'arm'an-Prandtl law should be replaced, when the viscosity is small but finite, by a Reynolds-number-dependent power law. In the case of local structure, an analogous argument shows that the Kolmogorov scaling of the second and third order structure functions is exact in the limit of vanishing viscosity, when the turbulence is most intermittent and least organized. When the viscosity is non-zero (Reynolds number large but finite), Reynolds-number-dependent corrections to the Kolmogorov-Obukhov scaling of the structure functions appear, due to a viscosity-induced reduction in intermittency. For higher-order structure functions the vanishing viscosity limit ceases to exist because of intermittency, and the Kolmogorov-Obukhov scaling fails. The near-equilibrium statistical theory we shall present is the basis of vanishing-viscosity asymptotics and relates the the behavior of the higher-order structure functions to the presence of intermittency. All parts of our analysis are heterodox in the context of the current state of turbulence research, but not in the broader context of the statistical mechanics of irreversible phenomena. 2. The intermediate region in wall-bounded turbulence Consider wall-bounded turbulence, in particular fully developed turbulence in the working section, i.e., far from the inlet and outlet, of a long cylindrical pipe with circular cross-section. It is customary to represent u, the time-averaged or ensemble-averaged longitudinal velocity in a pipe in the dimensionless form where u is the "friction" velocity that defines the velocity scale: where ae is the density of the fluid and - is the shear stress at the pipe's wall, d Here \Deltap is the pressure drop over the working section of the pipe, L is the length of the working section, and d is the pipe's diameter. The dimensionless distance from the pipe wall is where y is the actual distance from the wall and - is the fluid's kinematic viscosity. The length scale -=u in (2.4) is typically very small - less than tens of microns in some of the data discussed below. The key dimensionless parameter in the problem is the Reynolds number ud where - u is the velocity averaged over the cross-section. When the Reynolds number Re is large, one observes that the cross-section is divided into three parts (Figure 1): (1) the viscous sublayer near the wall, where the velocity gradient is so large that the shear stress due to molecular viscosity is comparable to the turbulent shear stress, a cylinder (2) surrounding the pipe's axis the where the velocity gradient is small and the average velocity is close to its maximum, and the intermediate region (3) which occupies most of the cross-section and on which we shall focus. The velocity gradient @ y u, (@ y j @ @y ), in the intermediate region (3) of Figure 1 depends on the following variables: the coordinate y, the shear stress at the wall - , the pipe diameter d, the fluid's kinematic viscosity - and its density ae. We consider the velocity gradient @ y u rather than u itself because the values of u depend on the flow in the viscous sublayer where the assumptions we shall make are not valid. Thus Dimensional analysis gives y ud where \Phi is a dimensionless function, or, equivalently, Outside the viscous sublayer j is large, of the order of several tens and more; in the kind of turbulent flow we consider the Reynolds number Re is also large, at least 10 4 . If one assumes that that for such large values of j and Re the function \Phi no longer varies with its arguments and can be replaced by its limiting value \Phi(1; (this is an assumption of "complete similarity" in both arguments, see [1]), then equation (2.8) yields and then an integration yields the von K'arm'an- Prandtl "universal" logarithmic law where - ("von K'arm'an's constant'') and B are assumed to be ``universal'', i.e. Re independent constants. The assumption that B is Reynolds-number-independent is an additional assumption. The resulting law is widely used to describe the mean velocity in the intermediate region; the values of - in the literature range between :36 and :44 and the values of B between 5 and 6.3- an uncomfortably wide spread if one believes in the "universality" of (2.10). However, there is no overwhelming reason to assume that the function \Phi has a constant non-zero limit as its arguments tend to infinity, nor that the integration constant remains bounded as Re tends to infinity. When either assumption fails other conclusions must be reached. Rather than list alternative assumptions we present a model problem that exhibits in a simple manner what goes wrong as well as the cure. Consider the equation dy y for y positive, where u is subject to the boundary condition positive parameter. One can view ffi as a dimensionless viscosity, and thus ffi \Gamma1 is analogous to a Reynolds number. One could reason as follows: For dy is approximately zero, and thus u is a constant, which can only be the constant 1. We can derive the same result for small y and ffi by an assumption of complete similarity: Equation (2.11) is homogeneous in the dimensions of u and y, and thus one can view both of these variables as dimensionless. By construction, ffi must be dimensionless. The dimensionless relation between these variables takes the form and if one makes an assumption of complete similarity, i.e., assumes that for ffi; y small \Phi is constant, one finds again that u is the constant 1. This conclusion is false. Equation (2.11) has the following solution that satisfies the boundary condition Note that for any positive value of ffi this solution is a power law and not a constant. We can obtain this solution for small y and ffi by assuming that the the solution is a power of the variable y while the form of its dependence on ffi is unknown; this leads to . (This is an assumption of "incomplete similarity" in y and no similarity in ffi). A substitution into equation (2.11) yields A(ffi); ff(ffi). Further, consider the solution (2.13) and, for any non-zero value of y, its limit as Clearly, and thus, as e, i.e., the limit of (2.13) for y ? 0 is the constant e. As deduced from the false assumption of complete similarity, the limit of u is a constant, but it is not the obvious constant 1. Furthermore, for a finite value of ffi, however small, u is not uniformly constant; it is not equal to e either for y ! ffi or for y large enough. The approximate equality u - e holds, when ffi is small but finite, only in an intermediate range constitutes an example of "intermediate asymptotics" [1]. Now consider subjecting \Phi in (2.8) to an analogous assumption, of incomplete similarity in j and no similarity in Re when both arguments are large [2]: Note a clear difference between complete and incomplete similarity. In the first case the experimental data should cluster in the (ln j; OE) plane on the single straight line of the logarithmic law; in the second case the experimental points would occupy an area in this plane. Both similarity assumptions are very specific. The possibility that \Phi has no non-zero limit yet cannot be represented asymptotically as a power of j has not been excluded. Both assumptions must be subjected to careful scrutiny. In the absence of reliable, high-Re numerical solutions of the Navier-Stokes equation and of an appropriate rigorous theory, this scrutiny must be based on comparison with experimental data. We now specify the conditions under which (2.15) may hold and narrow down the choices of A(Re) and ff(Re) (see [4,5,6,7,9,10]). Fully developed turbulence is not a single, well-defined state with properties independent of Re; there may be such a single state in the limit of infinite Reynolds number, but experiment, even in the largest facilities, shows that fully developed turbulence still exhibits a perceptible dependence on Re. Fully developed turbulence has mean properties (for example, parameters such as A and ff in (2.15)) that vary with Reynolds number like K 0 are constants and " is a small parameter that tends to zero as Re tends to infinity, small enough so that its higher powers are negligible, yet not so small that its effects are imperceptible in situations of practical interest; the latter condition rules out choices such as We expect A(Re) and ff(Re) in (2.15) to have the form are universal constants; we have implicitly used a principle derived from the statistical theory of section 4, according to which the average gradient of the velocity profile has a well-defined limit as the viscosity - tends to zero [5,6,10]. This is the the vanishing-viscosity principle. We expand A(Re); ff(Re) in powers of " and keep the first two terms; the result is: Substitution of (2.17) into (2.15) yields For this expression to have a finite limit as - tends to zero one needs ff must tend to zero as Re tends to infinity like ( 1 Re ) or faster. The assumption of incomplete similarity, experiment, and the vanishing-viscosity principle show that the threshold value Re is the proper choice. Use of this choice in equation (2.18) and an integration yield Re where the additional condition motivated by the experimental data, has been used. According to this derivation, the coefficients C are universal constants, the same in all experiments of sufficiently high quality performed in pipe flows at large Reynolds numbers [2,3]. In [12] the proposed law for smooth walls (2.19) was compared with the data of Nikuradze [26] from Prandtl's institute in G-ottingen. The comparison yielded the coefficients ff with an error of less than 1%. The result is We now wish to use the law (2.20) to understand what happens at larger Reynolds numbers and for a broader range of values of j than were present in the experiments reported by Nikuradze. If this extrapolation agrees with experiment, we can conclude that the law has predictive powers and provides a faithful representation of the intermediate region. We have already stated that the limit that must exist for descriptions of the mean gradient in turbulent flow is the vanishing-viscosity limit, and thus one should be able to extrapolate the law (2.20) to ever smaller viscosities -. This is different from simply increasing the Reynolds number, as - affects j and - u as well as Re. The decrease in the viscosity corresponds also to what is done in the experiments: If one stands at a fixed distance from the wall, in a specific pipe with a given pressure gradient, one is not free to vary ud=- and independently because the viscosity - appears in both, and if - is decreased by the experimenter, the two quantities will increase in a self-consistent way and - u will vary as well. As one decreases the viscosity, one considers flows at ever larger j at ever larger Re; the ratio 3 Re tends to 3=2 because - appears in the same way in both numerator and denominator. Consider the combination 3 ln j=2 ln Re in the form =\Theta d According to [3], at small -, i.e. large Re, - u=u - ln Re, so that the term ln(-u=u ) in the denominator of the right-hand side of (2.21) is asymptotically small, of the order of Re, and can be neglected at large Re; because the viscosity - is small, the first term in both the numerator and denominator of (2.21) is dominant, as long as the ratio y=d remains bounded from below, for example by a predetermined fraction. Thus, away from a neighborhood of the wall, the ratio 3 ln j=2 ln Re is close to 3/2 (y is obviously bounded by d=2). Therefore is a small parameter as long as y ? \Delta, where \Delta is an appropriate fraction of d. The quantity exp(3 ln j=2 ln Re) is approximately equal to exp Re '- Re '- Re \Gamma2 According to (2.20) we have also exp and the approximation (2.22) can also be used in (2.23). Thus in the range of y where y ? \Delta, but y ! d=2, we find, up to terms that vanish as the viscosity tends to zero, and Equation (2.25) is the asymptotic slope condition: As - ! 0 the slope of the power law tends to a finite limit, the limiting slope, which is the right hand side of (2.25). The von K'arm'an-Prandtl law also subsumes an asymptotic slope condition, with a limiting slope 1=-; the limiting slope in equation (2.25), approximately larger than a generally accepted value for - \Gamma1 . One can view equation (2.24) as an asymptotic version of the classical logarithmic law, but with an additive term that diverges as the Reynolds number tends to infinity, and of course a different slope. The family of curves parametrized by Re has an envelope whose equation tends to with close to the values of - found in the literature. The corresponding value of 1 - is exactly e times smaller than the value on the right-hand side of (2.25). It is clear that the logarithmic law usually found in the literature corresponds to this envelope; indeed, if one plots points that correspond to many values of Re on a single graph (as is natural if one happens to believe the von K'arm'an-Prandtl law (2.10)), then one becomes aware of the envelope. The visual impact of the envelope is magnified by the fact that the small y part of the graph, where the envelope touches the individual curves, is stretched out by the effect of - on the values of ln j. Also, measurements at very small values of y where the difference between the power law and the envelope could also be noticeable are missing because of experimental difficulties very near the wall. Thus, if our proposed power law is valid, the conventional logarithmic law is an illusion which substitutes the envelope of the family of curves for the curves themselves. The discrepancy of e between the slope of the curves and slope of the envelope is the signature of the power law, and helps to decide whether the power law is valid. The situation is summarized in Figure 2, which shows schematically the individual curves of the power law, their envelope, and the asymptotic slope. Historically, the understanding of the flow in the intermediate region of wall-bounded turbulence has been influenced by the "overlap" argument of Izakson, Millikan and von Mises (IMM) (see e.g [25]). This argument in its original form contains an implicit assumption of complete similarity, and once freed from it yields yet again the asymptotic slope condition (2.25). For details, see [5,7,10]. Detailed comparisons of the the power law and the von K'arm'an-Prandtl laws with experimental data are available in refs. [9,10]. For completeness, we exhibit in Figure 3 a set of experimental curves from the Princeton superpipe experiment [34]. Note its qualitative similarity to Figure 2. In particular, despite a flaw discussed in detail in [10], these experiments do indeed exhibit a separate curve for each Reynolds number, and a well-defined angle between the curves and their envelope. The applicability of the our analysis of the intermediate region of pipe flow to the velocity profile in a zero-pressure-gradient boundary layer is discussed in [11]. 3. Local structure in turbulence The analogy between the inertial range in the local structure of developed turbulence and the intermediate range in turbulent shear flow near a wall has been noted long ago (see e.g. [14,32], and it motivates the extension of the scaling analysis above to the case of local structure, where the experimental data are much poorer. In the problem of local structure the quantities of interest are the moments of the relative velocity field, in particular the second order tensor with components difference between x and x + r. In locally isotropic incompressible flow all the components of this tensor are determined if one knows is the velocity component along the vector r. To derive an expression for DLL assume, following Kolmogorov, that for it depends on h"i, the mean rate of energy dissipation per unit volume, r, the distance between the points at which the velocity is measured, a length scale , for example the Taylor macroscale T , , and the kinematic viscosity -: where the function f should be the same for all developed turbulent flows. If r is large, other variables may appear, as a consequence of external forces or of boundary conditions. The most interesting and the most important argument in this list is the rate of energy dissipation ". Introduce the Kolmogorov scale K , which marks the lower bound of the "inertial" range of scales in which energy dissipation is negligible: Clearly, the velocity scale appropriate to the inertial range is and this yields a Reynolds number The inertial range of scales is intermediate between the scales on which the fluid is stirred and the scales where viscosity dissipates energy, and is the analog of the intermediate region in wall-bounded flow. In this range the scaling law that corresponds to (2.15) is: r where as before, the function \Phi is a dimensionless function of its arguments, which have been chosen so that under the circumstances of interest here they are both large. If one now subjects (3.6) to an assumption of complete similarity in both its arguments, one obtains the classical Kolmogorov 2=3 law [21] from which the Kolmogorov-Obukhov "5/3" spectrum [27] can be obtained via Fourier transform. If one makes the assumption of incomplete similarity in r= K and no similarity in Re, as in the case of wall-bounded flow, the result is r where C; ff are functions of Re only. As before, expand C and ff in powers of 1 Re and keep the two leading terms; this yields Re r (ff 0 has been set equal to zero so that DLL have a finite limit as - ! 0). In real measurements for finite but accessibly large Re, ff small in comparison with 2/3, and the deviation in the power of r in (3.9) could be unnoticeable. On the other hand, the variations in the "Kolmogorov constant" have been repeatedly noticed (see [25,29,31]). Complete similarity is possible only if A 0 6= 0, when one has a well-defined turbulent state with a 2/3 law in the limit of vanishing viscosity, and finite Re effects can be obtained by expansion about that limiting state. In the limit of vanishing viscosity there are no corrections to the "K-41" scaling if equation (3.9) holds; this conclusion was reached in [15] by the statistical mechanics argument summarized in section 4 below. Kolmogorov [21] proposed similarity relations also for the higher order structure functions repeated p times; the scaling gives D Experiments in a small wind-tunnel by Benzi et al. [13], show some self-similarity in these higher-order functions, obviously incomplete, so that D LL:::L is proportional to r ip , with exponents i p always smaller then p=3 for p - 3, so that i instead of 1:67, i instead of 2:33, and i of 2:67. It is tempting to try for an explanation of the same kind as for 2: Re in other words, to assume that at Re = 1 the classic "K41" theory is valid, but the experiments are performed at Reynolds numbers too small to reveal the approach to complete similarity. If this explanation were correct, the coefficients ff p would be negative starting with 4, where there would be a reversal in the effect of the Kolmogorov scale (or whatever scale is used to scale the first argument in \Phi). As is well-known, for the Kolmogorov scaling is valid with no corrections. For must proceed with caution. We would like to present a simple argument that casts doubt on the good behavior of the structure functions for in the vanishing- viscosity limit. As Re ! 1 the intense vorticity and the large velocities in the fluid become concentrated in an ever smaller volume [15]. This is what we call "intermittency". is the fraction of the volume of a unit mass of fluid where the kinetic energy - u 2 is large, then ; one can see that fourth moments such as hu 4 i diverge as This casts a strong doubt on the good behavior of the fourth-order structure functions as the viscosity tends to zero; in the absence of such good behavior our expansions in powers of 1 Re cannot be justified and the explanation of the experimental data must proceed along different lines. Note that is the power where the sign of the power of r in an expansion in powers of 1 Re would change. The analysis just given of the second-order structure function contradicts the conclusions of the Benzi et al. [13], according to whom the asymptotic exponent in (3.9) is independent of Re and different from 2=3. We wish to point out however that, as we understand the discussion in [13], the exponent was found to be different from 2=3 only once it was assumed that it was not dependent on Re; to the contrary, even a cursory view of Fig. 3 in ref. [13] shows a marked dependence on Re. We are looking forward to an opportunity to reexamine these data in the light of our hypotheses. There is a key difference between a derivation of the Kolmogorov-Obukhov exponent from an assumption of complete similarity and its derivation as the vanishing-viscosity limit of an expression derived from an assumption of incomplete similarity. Complete similarity typically holds in statistical problems that are well-described by mean-field theories, while incomplete similarity typically applies to problems where fluctuations are significant. This remark is consistent with our conclusion, presented in [7,10], that the Kolmogorov scaling already allows for intermittency, and that its application to higher-order structure functions is limited by this very intermittency. 4. A near-equilibrium theory of turbulence At large Reynolds numbers Re the solutions of the Navier-Stokes equations are chaotic, and the slightest perturbation alters them greatly. The proper object of a theory of turbulence is the study of ensembles of solutions, i.e.,of collections of solutions with probability distributions that describe the frequency of their occurrence. We now outline a near-equilibrium theory of ensembles of flows on those small scales where the scaling theory of the previous section applies. This theory justifies the use of vanishing-viscosity asymptotics for appropriate moments of the velocity field and of its derivatives and supports the conclusions of the previous section regarding the behavior of the higher-order moments and structure functions. It is equivalent to earlier near-equilibrium theories [15], but the specific approach and the presentation are new; fuller detail can be found in [8]. We describe turbulence in terms of a suitable statistical equilibrium. In statistical mechanics, statistical equilibrium is what one finds in an isolated system if one waits long enough. One way of characterizing this equilibrium is by assuming that all states of the system compatible with the system's given energy can occur with equal probabilities; this is the "microcanonical ensemble". In turbulence the appropriate energy is the kinetic energy of the flow. An equivalent characterization is in terms of the "canonical ensemble", in which the probability of a state is proportional to exp(\GammafiH), where H is the energy of the state and fi is a parameter. In the canonical ensemble the energy is not fixed, and one can view the ensemble as describing a portion of an isolated system at equilibrium as it interacts with the rest of the system. The two ensembles are equivalent in the sense that, with a proper choice of the parameter fi and in a system with enough degrees of freedom, averages calculated in either ensemble are close to each other. The parameter fi is generally called the "inverse temperature" of the system. In many physical systems indeed proportional to what one intuitively perceives to be the temperature, as it can be gauged by touching a system with one's finger. However, the parameter fi can be viewed more abstractly, as the parameter that makes the two ensembles equivalent; in incompressible turbulence, in which there is no interaction between the macroscopic flow and the microscopic motion of the molecules of the fluid, the fi that one obtains cannot be gauged by the sense of touch. In a given system, fi is a function of the energy E and of whatever other variables are needed to describe the system. Note than in other realms of physics, for example in the kinetic theory of gases, one is well used to relating temperature to a suitable kinetic energy. Turbulence as a whole is generally not an equilibrium phenomenon: For example, if one stirs a box full of fluid and then isolates the resulting flow, the outcome after a long time is not turbulence in a statistical equilibrium but a state of rest; an isolated turbulent flow is one without outside forces or an imposed shear to keep it flowing. However, on the small scales in which we are interested, the relevant question is whether the motion has enough time to settle to an approximation of a statistical equilibrium in which one can assume that all the states with a given kinetic energy are equally likely to appear. The small scales have enough time when their characteristic time (length/velocity) is short enough compared to the characteristic time of the large-scale motion. An inspection of the Kolmogorov scaling given in the previous section shows that the characteristic time of an eddy of size r is proportional to r 2=3 , and small enough eddies (i.e., vortices) do have enough time to settle down to an equilibrium distribution. The task at hand is to construct the statistical equilibrium appropriate for turbulence, in particular specify its states. The question of how then to perturb it so as to take into account the irreversible aspects of turbulence has been treated elsewhere [16,17] and will not concern us here. Note that in most of the turbulence literature one speaks of the small scales reaching "equilibrium" when the energy distribution among them approximates the Kolmogorov-Obukhov form; here, for the moment, we mean by "equilibrium" a statistical equilibrium, in which all states have equal probability; we shall shortly claim that these two meanings are in fact identical. This is of course possible only if at the statistical equilibrium there are more states with much of the energy in the larger scales than states with much of the energy in the smaller scales. To agree with observation, a a hydrodynamical statistical equilibrium must have a finite energy density in physical space. To construct an equilibrium with this property, we start as in the construction employed in ref. [24] for vorticity fields; for simplicity we describe it in two space dimensions. Consider a unit box boundary conditions. Let be the velocity field and / a stream function; divide the box into N 2 squares of side h, in each square define a value of a discrete stream describe the location of the square; then define a discrete velocity field u ij by one-sided difference quotients of /, so that that the velocity is divergence-free. (In three dimensions there is one more index and the stream-function is replaced by a vector potential). The parameter h is an artificial cut-off, and we now present a procedure for letting this cut-off tend to zero while producing sensible fluid mechanics in the limit. Replace the energy R square dxdy by its discrete counterpart . For a fixed value of h, pick a value of E h , and, as a first step, assume that the values of u are equidistributed among all states with use a microcanonical ensemble. One can check that on the average, each one of the boxes has the same energy ju ij One may think that if one lets h ! 0 while keeping E h constant, the limit is an ensemble with a finite energy per unit volume; this is not so. The sequence of ensembles one obtains as reasonable limit: As h ! 0 the number of degrees of freedom tends to infinity, and there is no sensible way to divide a finite energy equally among an infinite number of degrees of freedom; indeed, if the energy per degree of freedom is zero the limiting ensemble has zero energy and no motion, and if the energy per degree of freedom is positive the limiting ensemble has an infinite energy (for a more thorough mathematical discussion, see [6,8]). One can also see that the limit of these microcanonical ensembles is meaningless by considering the corresponding canonical ensembles: One can check that as the parameter fi in the sequence of canonical ensembles tends to infinity; one can show that the only ensembles with infinite fi have either no energy or an infinite energy. To find a way out of this dilemma one must modify these ensembles as h ! 0 so as to ensure that the limit exists. We do so by looking at what happens to the parameter fi and keeping it bounded; furthermore, we do so on the computer. This is the key point: To obtain a sensible continuum limit, we keep fi bounded by keeping the energy from becoming equally distributed among the degrees of freedom, and this produces an average energy distribution among scales that agrees with the Kolmogorov-Obukhov law and produces the Kolmogorov scaling of the low-order structure functions. One can also show (see [8]) that this very same procedure is needed to produce ensembles whose members, the individual velocity fields, do not violate what is known about the solutions of the Navier-Stokes or Euler equations. To proceed, we have to be able to calculate fi given h and E h . Averages with respect to microcanonical ensembles can be calculated numerically by an algorithm known as "mi- crocanonical sampling" [18]: Introduce an additional variable, a "demon", which interacts with all the degrees of freedom in some random order. In each interaction, the demon either absorbs an energy packet of some predetermined magnitude s, s !! E h , or gives away an energy packet of the same size. If the demon takes in an energy s, it reduces the energy in the velocity field by modifying / ij so that the integral R dxdy over the unit square is reduced by s; the effect of this reduction modifies the values of u in the neighborhood of the point (i; j). If the demon gives out energy, it modifies / so as to increase the energy integral. The demon is constrained so that it cannot give out energy unless it had acquired energy in its previous history; no "loans" are allowed. The sequence of states wrought by the demon's actions ranges over even-handedly the configurations of the system. If one wishes to conserve an additional quantity, as we shall below, one can do so by allowing the demon to exchange doses of the conserved quantity as it wanders along, subject to the condition that it never give out what it does not have. The parameter fi in the equivalent canonical ensemble can be determined in the course of calculating averages: As the demon interacts with the ensemble it typically has some energy stored away; the system consisting of the physical system plus the demon is isolated, and by the equivalence of the canonical and microcanonical ensembles, the probability of an energy E d being stored by the demon is canonical, i.e., proportional to exp(\GammafiE d ); this observation allows one to estimate fi after the demon has had a sufficient number of interactions. In addition to its dynamical role in moving the system from state to state for the purpose of calculating averages, the demon reveals the value of fi; if h and E h are given, there is a well-defined numerical procedure for finding the corresponding fi. Rather than keep fi merely bounded, we keep it constant. To do this, one needs a variable that can be altered and whose variation controls fi. Experience and mathematics show that one can use as control variable the integral I = R j-jdxdy , where - is the vorticity calculated by finite differences and the integral is approximated by the appropriate sum. Thus the plan is to determine an I for each h so as to keep fi at a fixed value fi goal common to all the h. For simplicity and without loss of generality we set E For a given h, pick a starting guess for I, say I 0 , and then produce a sequence of better values I n by the formula where fi without a subscript is the latest estimate of fi available from the demon and K is a numerical parameter chosen so as to ensure that the I n converge to a limit. Before calculating a new value I n+1 of I the demon must be allowed at least one energy exchange with the ensemble, during which the variable I is maintained at its last value I n . Once fi reaches the desired value fi goal the quantity I remains constant. The resulting ensemble gives non-zero, equal probabilities to all states compatible with both the given value and the calculated value of I; when both constraints are satisfied, the energy per degree of freedom is no longer the same for all the degrees of freedom. The changes in I needed to keep fi fixed as h is changing are displayed in Figure 4 for several values of . The statistical error is throughout of the order of 2%. The values of I needed to keep fi fixed increase with As shown in [8, 16], for small enough h the curves are independent of the value of fi, and this fact is reflected in the confluence of the several curves in Figure 4. Note that I is calculated on the grid by taking differences of the values of the velocity u at points separated by h, which by the Kolmogorov-Obukhov scaling should be proportional to h 1=3 ; then one divides by h, and takes an average; one expects I to grow with N like N 2=3 . In Figure 5 we plot the logarithm of I vs. the logarithm of N ; the relation is well approximated by a straight line whose slope is :65 with an error of \Sigma:05. Within the limitations of the Monte-Carlo sampling, the Kolmogorov-Obukhov scaling is seen to be applicable in this equilibrium model. As the Kolmogorov-Obukhov scaling applies to the low-order structure functions in a flow with a finite but small viscosity, Figure 5 shows that low-order moments structure function have a limit as the viscosity tends to zero. One can perform a similar analysis of the small-scale structure of flow near a wall and conclude that the the first-order moments of the derivatives of the velocity field near walls have well-behaved limits, a fact used above in the discussion of the scaling of the wall-region in the pipe. Figure 4 defines the limiting process in which h ! 0 with a limit that provides a meaningful equilibrium ensemble for the small scales of the flow. The fact that the construction above is numerical enhances its value rather than detracts from it, as we expect to use similar constructions in the numerical modeling of turbulence. An elegant argument, suggested by the work of Kailasnath et al. [19] and presented in detail in [8], shows that the third-order structure function is calculated in the equilibrium theory exactly; of greater interest here is what happens to moments of the velocity field of order four and more. We have argued in the preceding section that the vanishing viscosity limit is not well-behaved for these higher moments, and thus the good behavior of the structure functions is unlikely and the expansion in powers of 1= ln Re is invalid. In the equilibrium theory the fourth-order moments fail to converge to a finite limit as In Figure 6 we display the fourth-order moment R juj 4 dxdy as a function of N at the parameter value 5. Higher-order moments diverge even faster. The divergence of the higher moments corresponds to the formation of concentrated vortical structures, like the ones explicitly constructed in [16]. We have thus produced Kolmogorov scaling for the low-order moments in a system that is highly intermittent in the sense that the vorticity is concentrated on a small fraction of the available volume. The results of the equilibrium theory are therefore consistent with the scaling analysis of the previous section, according to which the Kolmogorov scaling of the second-order structure function is exact in the limit of vanishing viscosity not despite intermittency but because of intermittency, while its failure for the higher-order moments can be ascribed to the absence of a well-behaved vanishing-viscosity limit, as a result of which the expansion in the inverse powers of ln Re is not legitimate. Note the small number of assumptions made in the equilibrium theory; all that was assumed was that the fluid was near statistical equilibrium on the small scales, the fluid was incompressible, the energy density in physical space was finite, and a probability measure on the ensemble of flows was well-defined. The Navier-Stokes equations did not enter the argument in the present paper (but see ref. [8]). Finally, it is worth noting that an analysis of simplified near-equilibrium vortex models [22,33] has provided an example where an expansion in powers of a parameter analogous to 1 ln Re can be fully justified without recourse to experimental data. 5. Conclusions We have reached the following conclusions: (1) The von K'arm'an-Prandtl law of the wall must be jettisoned, and replaced by a power law with a Reynolds-number dependent coefficient and exponent, as suggested by an assumption of incomplete similarity. (2) The Kolmogorov-Obukhov scaling of low-order structure functions in the local structure of turbulence admits only viscosity-dependent corrections, which vanish as the Reynolds number tends to infinity. There are no "intermittency corrections" to this scaling in the limit of vanishing viscosity. The Kolmogorov scaling of the higher-order structure functions fails because of intermittency. (3) These conclusions are consistent with and support the near-equilibrium theory of turbulence. Acknowledgement Prof. Barenblatt and I would like to thank the following persons for helpful discussions and comments and/or for permission to use their data: Prof. N. Gold- enfeld, Prof. O. Hald, Dr. M. Hites, Prof. F. Hussain, Dr. A. Kast, Dr. R. Kupferman, Prof. H. Nagib, Prof. C. Wark and Dr. M. Zagarola. --R Cambridge University Press Reynolds number) in the developed turbulent flow in pipes Basic hypotheses and analysis of local structure and for the wall region of wall-bounded turbulence turbulence theory asymptotics and intermittency in preparation Discussion of experimental data Sciences USA the zero-pressure-gradient turbulent boundary layer Part 2. Physica D 80 Applications to Physical Systems law and odd moments of the velocity difference in turbulence USSR Series 4 exponents at very high Reynolds numbers for the two-dimensional XY model Reynolds number Figure 1. Figure 2: Schematic of the power law curves The individual curves of the power law The envelope of the family of power law curves (often mistaken for a logarithmic The asymptotic slope of the power law curves. Figure 3. of the experimental data according to their Reynolds numbers and the rise of the curves above their envelope in the (ln j at the center of the pipe. and are incompatible with the von K'arm'an-Pradtl universal logarithmic law Figure 4. Figure 5. given by the Kolmogorov-Obukhov scaling Figure 6. --TR

wall-bounded turbulence;local structure;intermittency;statistical theory;turbulence;scaling

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Boundary representation deformation in parametric solid modeling.

One of the major unsolved problems in parametric solid modeling is a robust update (regeneration) of the solid's boundary representation, given a specified change in the solid's parameter values. The fundamental difficulty lies in determining the mapping between boundary representations for solids in the same parametric family. Several heuristic approaches have been proposed for dealing with this problem, but the formal properties of such mappings are not well understood. We propose a formal definition for boundary representation. (BR-)deformation for solids in the same parametric family, based on the assumption of continuity: small changes in solid parameter values should result in small changes in the solid's boundary reprentation, which may include local collapses of cells in the boundary representation. The necessary conditions that must be satisfied by any BR-deforming mappings between boundary representations are powerful enough to identify invalid updates in many (but not all) practical situations, and the algorithms to check them are simple. Our formulation provides a formal criterion for the recently proposed heuristic approaches to persistent naming, and explains the difficulties in devising sufficient tests for BR-deformation encountered in practice. Finally our methods are also applicable to more general cellular models of pointsets and should be useful in developing universal standards in parametric modeling.

consumer applications, probably has to do more with development and successful marketing of new parametric ("feature-based" and "constraint-based") user interfaces than with the mathematical soundness of solid modeling systems. These parametric interfaces allow the user to define and modify solid models in terms of high-level parametric definitions that are constructed to have intuitive and appealing meaning to the user and/or application (see [16] for examples and references). The success of parametric solid modeling came at a high price: the new solid modeling systems no longer guarantee that the parametric models are valid or unambiguous, and the results of modeling operations are not always predictable. Unpredictable behavior of the new systems is well documented in the literature [11, 13, 15, 34], and we will consider some simple illustrative examples in the next section. The basic technical problem is that a parametric solid model corresponds to a class of solids, but there is no formal definition or standard for what Complete address: 1513 University Avenue, Madison, Wisconsin 53706, USA. Email: [email protected], ParameterInstance 'persistent naming' ParameterInstance Figure 1: The role of persistent naming for a parametric family of solids this class is [34]. As a result, different modeling systems employ incompatible, ad hoc, and often internally inconsistent semantics for processing parametric models. The lack of formal semantics manifests itself clearly through the so called "persistent naming" problem [15, 34]. In modern systems, every edit of a parametric definition is followed by a regeneration (boundary evaluation) of the boundary representation (b-rep) for the resulting solid. Since parametric definitions and constraints explicitly refer to entities (faces, edges, vertices) in the boundary representation, every such regeneration must also establish the correspondence between the pre-edit and post-edit boundary represen- tations. The correspondence must persist over all valid edits, hence the term "persistent naming". But since the semantics of a parametric family is not well-defined, neither is the required correspondence. All systems seek to establish the correspondence that is consistent with the user's intuition, but no reliable methods for achieving this goal are known today. The problem is illustrated in Figure 1. Every two instances of parametric definition are related through a parametric edit of one or more parameters t; each instance has its own b-rep(t), and the problem of persistent naming is to construct the map g between the two corresponding b-reps. Recently, several schemes for persistent naming have been proposed [10, 18, 25, 31] that appear to work well in a variety of situations. But the proposed techniques are not always consistent between themselves and none come with guarantees or with clearly stated limitations; they may be able to alleviate the observed problems in today's systems, but not solve them - because the problems have not even been formulated. 1.2 Main contribution The diagram in Figure 1 places no constraints on properties of the naming map g. A commonly cited obstacle to such formal properties stems from the belief that in the most general case parametric edits do not seem to preserve the basic structure of the boundary representation. For example, parametric edits may lead to collapsed and expanded entities, (dis)appearance of holes, etc. Yet the naming mechanisms proposed in [10, 18] clearly rely on the generic topological structure of the boundary representation as a primary naming mechanism. Certainly we would expect that the naming problem should have a unique solution when the parametric changes are small and the resulting boundary representation does not change very much. Accordingly, we propose that any notion of a parametric family should be based on the following principle of continuity : ffl Throughout every valid parameter range, small changes in a solid's parameter values result in small changes in the solid's representation. This statement will be made precise in Section 2 of the paper. We argue that this essential requirement of continuity eliminates much of the unpredictable and erratic behavior of modern systems as described in [15, 34]. In this paper, we use this principle of continuity to develop a formal definition for Boundary Representation (BR-)deformation for solids in the same parametric family. Importantly, "small changes" may include limited local collapsing of entities in boundary representations; in this case the resulting BR-deformation can actually change the topology of the boundary representation. A special case of the topological BR-deformation is a geometric BR-deformation that may modify the shape of the boundary cells, but the topology and the combinatorial structure of the b-rep remains unchanged. Our formulation does not solve the difficult problem of constructing the naming map g in Figure 1, but it leads to necessary conditions that must be satisfied by any naming map g between boundary representations in the same parametric family. Loosely, the oriented boundary of each cell in the pre-edit b-rep must be mapped into the closure of the corresponding cell in the post-edit b-rep with the same orientation. These conditions are powerful enough to identify invalid updates in many (but not all) practical situations, and the algorithms to check them are simple. In particular, the identified conditions give formal justification to the naming mechanisms proposed in [10, 18], but also identify techniques that are not acceptable, because they do not satisfy the assumption of continuity. Finally, our definitions explain the difficulties of devising sufficient conditions for BR-deformation and the associated naming map g. A definition for BR-deformation is a prerequisite for extending informational completeness to parametric solid modeling and for creating an industry-wide standard for data exchange of parametric models. To our knowledge, there are no competing proposals. A notion of variational class is prominent in mechanical tolerancing and robustness [29, 36]. Notably, Stewart proposed a definition of variational class in [36], which was later used in study of polyhedral perturbations that preserve topological form in [3]. The assumptions of continuity and continuous deformation are also mentioned informally in the naming techniques and discussion of [11]. Our definitions are consistent with the earlier definitions, but are stronger and stated in algebraic topological terms, which are more appropriate for dealing with boundary representations and the mappings between them. The proposed notion of parametric family and the principle of continuity are stated in terms of a particular representation scheme (boundary representation in our case). But our formulation and results are applicable to more general cellular representation of pointsets. In Section 5 we discuss briefly how our approach to parametric modeling may be extended to deal with arbitrary parametric edits and other representations. Applying similar principles of continuity to other representations would lead to different definitions of parametric families, e.g. those based on small changes in CSG representations [30, 31, 34]. We do not consider such families in this paper because most extant commercial systems are based on b-reps. 1.3 Examples The following examples illustrate some of the erratic and inconsistent behavior that occurs during parametric updates of a solid model. These examples were created in several versions of two leading commercial solid modeling systems. 1 The simple solid S shown in Figure 2(a) is sufficient to illustrate many of the relevant issues. The shown boundary representation was produced from a parametric definition of S. The parametric definition uses a number of parameters, including parameters t and d that constrain the locations of features (slot and hole) with respect to the solid's edges, as shown in the figure. The value of t determines the location of which in turn determines the location of the hole constrained by parameter d. Consider what happens during a parametric edit of t. Every time t is changed, the system attempts to regenerate the boundary representation of the solid and establish which of the edges should be called . The result of this computation is easy to determine by simply observing what happens to the hole constrained by d, since it explicitly references edge e 1 If the change in t is small enough not to produce any changes in the combinatorial structure of the boundary representation, we would expect to see no abrupt changes resulting from the edit. In other words, the hole and the constraint d should move continuously with changes in the value of t. Instead, Figure 2(b) illustrates an example of situations that have been observed in earlier versions of some systems. In some cases the hole jumps on the face f 0and in other situations the hole surprisingly jumps to the other side of the edge e 0onto f 0. In either case, the observed behavior would indicate that e 1 was mapped to e 0; in the latter case the placement of the hole was also affected by the orientation of the mapped edge e 0(which is opposite to that of e 0with respect to face f 0). In both cases the assumption of continuity is violated, even though the correct continuous mapping clearly exists. We will formally characterize such a mapping in section 2.3 as geometrically deforming. Similar problems have been identified and illustrated with "jumping blends" by Hoffmann [10]. The naming techniques proposed in [10, 11, 18], as well as the latest versions of 1 We do not identify the specific systems, because similar technical problems are easily identified in all parametric modeling systems that rely on b-rep for internal representation of solids. (a) Parameter t controls the b-rep of the solid (b) Erratic behaviour observed in earlier systems (c) Continuous change in t collapses face f1 (d) B-rep is not in parametric family of (a) d d d d d Figure 2: Parametric edits of a simple solid many commercial systems, are now able to establish the appropriate mapping in this simple case, but not in many more general situations. Figure 2(c) shows another parametric update of the same solid, where the value of t is equal to the radius of the cylindrical slot, and the edge e 1 becomes coincident with e 3 . Few (if any) of the latest commercial systems can handle such an update in a consistent and uniform fashion. A typical response in one system is shown in Figure 2(d), which indicates again that e 1 was mapped to e 0; another system simply signaled an error indicating that it could not locate the proper edge (the naming map could not be found). When the slot was constructed as a "hole," the latter system signals no error and simply maps e 1 to e 0, with the smaller hole reappearing on the other side of e 0(outside the solid). Yet another system might delete the constraint (and the hole) with an appropriate warning message. Again, intuitively it is clear that continuous change in should lead to the collapse of face f 1 and identification of edges e 1 and e 3 to e 0. We will characterize such collapsing maps in section 2.3 as topologically deforming. These simple examples clearly show that the naming map is usually developed in ad hoc fashion, is dependent on the way the solid is created, and may be internally inconsistent across different features and representation schemes. The matching methods proposed in [10, 18] will not always properly handle collapsing maps, although it appears that the techniques in [10] could be further adapted. One may also start with the solid shown in Figure 2(d), i.e., create the hole and constrain it with respect to e 0, and then increase the value of parameter t. One would then expect the face f 0and the hole to gradually move to the right; we will call the corresponding mapping between pre-edit and post-edit boundary representations expanding. Unfortunately, a more typical response from a system is the one shown in Figure 2(a), and the result will certainly differ from version to version and from system to system. (a) Original solid's b-rep K (b) Modified solid's b-rep L Figure 3: Discontinuous update allowed by the parametric modeling system Our final example is shown in Figure 3. Parameter t controls the location of the cylindrical feature that is unioned to the cubical body, and one of the intersection edges is blended. 2 The change in the value of t moves the cylinder from one side of the block to the other. The system matches the blended edge e 1 with the edge e 0in the post-edit b-rep as shown. Is there a way to determine if this update is correct? Notice that there exists a continuous map between pre-edit and post-edit boundary representations: it is a simple rigid body motion that flips S upside-down. But this map is not consistent with the parametric change in t, which requires that face f 1 is mapped into face f 0. While some of the proposed naming techniques will allow such edits, we will see in section 3.4 that no change continuous in t could produce such a boundary representation. All boundary representations are cell complexes, and their formal properties, including validity, are rooted in algebraic topology [14, 27]. Algebraic topology is also the proper setting for formulating BR-deformation. Below we use a number of standard concepts (complexes, chains, continuous maps, etc.) that can be found in many textbooks [1, 5, 17, 22]. For brevity, we only state few critical definitions and results that are necessary for formulating and understanding the concept of BR-deformation. Blends provide a convenient way to track the names of the edges [10]. Throughout the paper we will ignore the geometric changes in b-rep resulting from blended edges. 2.1 Boundary representation as a cell complex The boundary of a 3-dimensional solid is an orientable, homogeneously 2-dimensional cell complex. Internally, every b-rep consists of an abstract (unembedded) cell complex, that captures all combinatorial relationships between cells (vertices, edges, and faces), and geometric information that specifies the embedding for every cell in E 3 . We do not deal with specific data structures in this paper, and therefore it is not necessary to distinguish between the abstract cell complex and its embedding. Thus, we will assume that the boundary of a solid S is represented by an embedded complex K (the b-rep of S). Many different cell complexes have been used for boundary representations, including simplicial, polyhe- dral, and CW (with cells homeomorphic to disks). Our discussion, definitions, and results are not affected by any particular choice of a cell complex. As the example in Figure 2 illustrates, b-reps of many commercial systems may contain more general cells, such as faces with holes, circular edges, and so on; such cell complexes are also natural from the user's point of view. The corresponding formal definition, which we also adopt in this work, is given by Rossignac and O'Connor in [23]: a Selective Geometric Complex (SGC) is composed of cells that are relatively open connected submanifolds of various dimension and are assembled together to satisfy the usual conditions of a cell complex: 1. The cells in K are all disjoint; and 2. The relative 3 boundary boe of each p-cell oe is a finite union of cells in K. Both SGCs and our definitions are valid in n-dimensions, but for the purposes of solid modeling, we will assume that each p-cell oe of K is a p-dimensional subset of E 3 with 3. The union of all embedded cells is the underlying space jKj of the complex K. In particular, the set of points b(S) of every solid S is the underlying space of some boundary representation K. The boundary of a homogeneously p-dimensional complex K can be also characterized algebraically, as an algebraic topological chain @C(K). The precise difference and the significance of this distinction will be become apparent in section 3. 2.2 Cell maps and cell homeomorphisms The cellular structure of b-reps implies the proper mechanism for defining the notion of a variational class: we will consider the effect of a parametric change on the b-rep cell by cell. Informally, if K is the pre-edit b- and L is the post-edit b-rep(t 1 ), then the naming map g seeks to establish a correspondence between cells in K and cells in L. So for example, when t 1 , we have and it is natural that g should be a cell-by-cell identity mapping. When the difference t 1 is so small that none of the cells collapse, the correspondence between the cells does not change, even though some cells may continuously change their shapes. For example, in Figure 2(a), small changes in t correspond to a small movement of the hole on face f 2 and therefore to a change in the shape of f 2 in the post-edit boundary representation. As the difference larger, so does the difference between the shapes of the corresponding cells; and eventually some cells may actually continuously collapse into some other cells. This is exactly the case in Figure 2(c) where face f 1 and all of its bounding edges collapsed into the edge e 3 and its bounding vertices. In other words, the cells in the post-edit b-rep L are images of continuous variations of the corresponding cells in the pre-edit b-rep K. This cell-by-cell continuous variation is captured by the notion of a continuous cell map (modified from [17]). and L be two b-reps. Map called a continuous cell map from K to L if: 1. for each cell oe 2 K, g(oe) is a cell in L; 2. dimension(g(oe)) - dimension(oe); 3. for every oe 2 K, g restricted to the closure of oe is continuous. 3 The relative boundary of a cell oe is defined as the set difference between the closure of oe and oe itself [23]. The first two conditions capture the notion of a cell-by-cell map. The third condition implies that g is a continuous map on every cell that can also be extended to the continuous map on the whole boundary of the solid [1, 22, 24]. By definition, continuity of g requires that g(closure(oe)) ae closure(g(oe)) for every cell oe [21]. But since K and L are finite, and jKj and jLj are closed and bounded, the stronger condition can be shown to hold [22, p.215]. In keeping with the common practice in topology literature, a cell map g is used to denote both a map between two b-reps and a map between their underlying spaces, depending on the context in which it is used. In the same spirit, we will at times abuse the notation and use oe to denote both the cell or its underlying space, again depending on the context. Thus, cell map g plays a dual role: on the one hand, it acts as a "naming" map and establishes the correspondence between the cells of two boundary representations; on the other hand it guarantees the continuity of variation in the underlying space (the solid's boundary). Before we go any further, it may be instructive to check whether naming maps can be constructed for the examples of the previous section that are also cell maps. Consider the naming map g between the two boundary representations of the solids in Figure 2(a) and 2(c) defined by: and so on, i.e. the rest of the cells in Figure 2(a) map to their implied images in Figure 2(c). It is easy to verify that g is a cell map, because it satisfies all three conditions of the definition given above: every cell in (a) is continuously mapped to a cell in L of the same or lower dimension (note that face f 1 is mapped to edge e 0) and the continuity condition is satisfied for every cell. In particular, note that the boundary set of the collapsed face f 1 is mapped into the closure of e 1 . Now consider the mapping of the same boundary representation in (a) to the one shown in Figure 2(d): and the rest of the cells in (a) map to their appropriate images in (d). It is easy to check that g is not a cell map; for example, the closure of face f 1 contains the edge e 1 , while e 0= g(e 1 ) is not in the closure of Referring to Figure 3, consider the mapping g given by: where n is the number of faces and m is the number of edges respectively in the boundary representation. This map g trivially satisfies all three required conditions and is therefore a cell map. However, there is also another cell map h between the two boundary representations: This cell map h maps the loop of edges around face f 1 into the loop of edges around f 2 and vice versa, and corresponds to a rigid body rotation of 180 degrees about the shown x-axis. Thus, given two boundary representations K and L there may be more than one valid cell map between them. For sufficiently small parameter changes, we expect that none of the cells in the boundary representation collapse and that there is also a one-to-one correspondence between cells in K and L. In this case, we can strengthen the definition of a cell map: bijective cell map called a cell homeomorphism between jKj and jLj [22]. Notice that when g is a cell homeomorphism, the second condition in the definition of the cell map can also be strengthened: the dimension of g(oe) must be equal to that of oe for every cell in K. For example, both cell maps g and h, constructed above for the transformation in Figure 3, are cell homeomorphisms. 2.3 Topological and geometric deformations In order for boundary representation L to be in the parametric family of boundary representation K, there must exist a (not necessarily unique) cell map g from K to L. Furthermore, the principle of continuity postulated in section 1.2 implies that when two boundary representations are in the same parametric range, one must be continuously deformable into the other, as illustrated in Figure 4. Unfortunately, the existence of a cell map from K to L alone is not sufficient to assure that K is deformable into L. (a) Figure 4: Continuous deformation from K to L For example, even though the two boundary representations in Figure 3 are related by a cell homeomorphism g, K cannot be deformed into L without leaving E 3 , because L is a reflection of K. We will see in section 4 that such cell maps can be characterized and detected in a straightforward fashion. A more difficult situation is illustrated in Figure 5. The boundary representation of two linked tori is related to the boundary representation of the two unlinked tori by a cell homeomorphism; yet it is not possible to deform one into another without leaving E 3 , because the continuity principle is violated for some intermediate values of the shown parameter t. If K is a b-rep(t 0 and L is a b-rep(t 1 ), then the continuity principle requires that, as t changes from t 0 to t 1 , K continuously deforms into L. These observations motivate the following definition. (a) Original cell complex K (b) Modified cell complex L Figure 5: K is homeomorphic to L, but K cannot be deformed into L without leaving E 3 Definition 3 (BR-deformation) Let K, L be two boundary representations. A map called a Boundary Representation deformation or BR-deformation of K to L if: 1. F (jKj; 2. F (jKj; 3. F (jKj; t) is a cell map that is also continuous in t 2 [0; 1]. Without loss of generality, we use symbol I in the above definition to denote the closed interval [0; 1], which corresponds to the normalized range of valid values for the parameter t. In standard algebraic topological terms, BR-deformation F is a homotopy [1, 17] and can be also viewed as a continuous family of cell maps F t , for every value t 2 I. In particular, F 0 is the identity cell map and F 1 is the cell (naming) map g from K to L. BR-deformation captures the spirit of the postulated requirement of continuity and appears to be the weakest possible condition. The obvious BR-deformation can be found by examination in Figure 4 for the parametric updates shown in Figures 2. Let us consider again the typical situation illustrated in Figure 1 with pre-edit b-rep(t 0 denoted by K and post-edit b-rep(t 1 denoted by L. Only the map g between the two boundary representations is known, and the problem is to determine whether an appropriate BR-deformation F exists or not. If BR-deformation does exists, then we say that g is a BR-deforming map. Definition 4 (BR-deforming map) A map g from boundary representation K to boundary representation L is called BR-deforming if F (jKj; BR-deformation F . If one accepts the principle of continuity and the implied definitions proposed above, then it may be reasonable to require that a naming map g for solids in the same parametric family must be BR-deforming. We constructed the appropriate BR-deformation in Figure 4 by examination, and it is easy to see that no BR-deformation exists for the example in Figure 5. But our definition of BR-deformation involves conditions not only on K and L but also on all the b-reps in the parametric range which are typically not known. Furthermore, such a BR-deformation need not be (and usually is not) unique. Is there a general method for checking if a given map g is BR-deforming? For example, it may not be clear why no appropriate BR-deformation exists for the parametric update shown in Figure 3. We will delay discussing the sufficient conditions for existence of BR-deformation until section 5. In the next section, we discuss necessary conditions for the existence of BR-deformation and show through examples how they can be used in practice to distinguish b-reps which belong to the same parametric family from b-reps which do not. Observe that boundary representations in Figures 4(a), (b) and (c) are related by a continuous family F t of cell homeomorphisms. This corresponds to a special but common and important case of geometric BR-deformation F . The combinatorial structure of the boundary representation K, including the dimension of all cells, is preserved under geometric BR-deformation. On the other hand, the further deformation of the solid in Figure 4(d) is not geometric BR-deformation, because one of the faces is continuously collapsed into the adjacent edge and the combinatorial structure of the resulting b-rep has changed. It should be clear that geometric BR-deformation is a special case of the more general topological BR-deformation. We shall see in the next section that the necessary conditions for geometric BR-deformation are simpler and are easier to test. Necessary Conditions for BR-Deformation 3.1 Combinatorial closure and star conditions The third condition in definition 1 of cell map g requires that g restricted to the closure of every cell is continuous. We already used this condition to test whether the maps for the examples of section 1 are cell maps. The closure of a cell oe in a complex K consists of oe itself and all lower dimensional cells incident on oe. Thus, the closure of a two-dimensional face includes its bounding edges (which are often represented by intersection of surfaces), and the closure of a vertex is the vertex itself. To facilitate development of more convenient algorithmic tests, we now restate this condition in alternative combinatorial terms, using the Cauchy's definition of continuity in terms of open neighborhoods [2]. The combinatorial equivalent of a cell's open neighborhood is the star of the cell. The following definition is modified from [9, 22] and is also consistent with definitions in [23]. Definition 5 (Star of a cell) The star of a cell oe in a complex K, denoted St(oe), is the union of oe and all cells in K that contain oe in their boundary. In other words, the star of a cell is the union of the cell with its neighboring higher-dimensional cells. For example the star of the edge e 1 of the solid shown in Figure 2(a) is the union of cells g. The star of a vertex is a union of the vertex with all its adjacent faces and edges. The star of a face is just the face itself, as there are no three-dimensional cells in boundary representations. We can now restate the requirement that g restricted to the closure of oe is continuous, originally expressed by the condition (1), to mean that every neighboring cell of oe is mapped to a neighboring cell of the image cell g(oe) or to g(oe) itself. Remembering that some cells oe may collapse into lower-dimensional cells g(oe), this requirement translates into a simple condition that must be satisfied by every cell[2, 22]: The strict inclusion applies when the cell oe is mapped to a lower dimensional cell, while condition (2) becomes an equality when the dimension of oe remains the same under the map g. In the latter case, this condition simply means that adjacency of all cells is also preserved under the map g. Let us consider the examples of the previous section to see how the star condition can be verified. The naming map between solids in Figure 2(a) and Figure 2(c) takes the star of e 1 into the union of their respective image cells fe 0; f 0g. The star of e 0is the union of cells fe 0; f 0g and the other (hidden) adjacent face of e 0. This clearly shows that g(St(e 1 )). The star condition (2) is easily verified to hold for other cells as well. Now consider the mapping of the same boundary representation in Figure 2(a) to the one shown in Figure 2(d). We already concluded earlier that this map is not a cell map because the closure condition is violated, and this can be easily verified by the star condition. The star of e 1 is the union of cells which are mapped to cells fe 0; e 0; f 0g. Since the star of g(e 1 ) is the union of fe 0; f 0; f 0g, it is clear that )), and g cannot be a cell map. 3.2 Oriented cell maps As we have seen above, the requirement on g to be a cell map is a necessary but very weak condition for BR-deformation. The difficulty with the example in Figure 3 is that the proposed cell map from K to L does not preserve the relative orientation between certain cells. A simpler two-dimensional example is shown in Figure 6. A cell map defined by g(oe 0 takes the cell complex (triangle) K into its reflected copy L; however it is not possible to continuously deform K into L without leaving E 2 . (a) Original cell complex K (b) Modified cell complex L f Figure Cell map between K and L does not preserve orientation In order to take advantage of the orientation information, we need to slightly modify the definitions of a cell complex and a cell map by requiring that all cells are oriented, which is usually the case in most boundary representations. Orientation of a cell oe 2 K can be visualized as a (positive or negative) sense of direction. More specifically, every 0-cell (vertex) can be oriented positively or negatively, orientation of every 1-cell (edge) bounded by two vertices s and t is determined by the order of s and t, and 2-cells (faces) can be oriented either counterclockwise (positive) or clockwise (negative). Cyclic edges that do not have vertices can be also directed in one of two ways. Oriented cell complexes are shown in Figures 3 and 6. We can now modify the definition of the cell map to account for orientation. Definition 6 (Oriented Cell Map) A map is an oriented cell map if it is a cell map and for every oriented cell oe 2 K, g(oe) is an oriented cell in L. It can be shown that orientation is a homotopy invariant and is automatically preserved through continuity [5], which means that every BR-deforming map is orientation preserving . In other words, the sense of direction must be preserved at every point of the solid's boundary, and we need to assure this condition in a cell-by-cell fashion. Since boundary representation of every solid is orientable [14, 27], the above condition is easily enforced on 2-cells by requiring that all faces are oriented positively (counterclockwise); we will also postulate that all 0-cells (vertices) have positive orientation. It is more difficult to state what it means to preserve the orientation on 1-cells; note however that orientation of a face induces orientation in its bounding edges, and an edge orientation induces orientation in its bounding vertices. For example, in Figure 6(a) the assumed orientation of edge e 2 is consistent with the orientation of e 2 induced by the face f . By contrast, in Figure 6(b) the orientation induced in e 0by f is opposite to the assumed orientation of e 0. A similar situation can be observed in Figure 3 for the edges bounding face f 3 . Additional discussion and examples of assumed and induced orientation can be found in most texts on algebraic topology, e.g. in [1]. 3.3 Induced orientation condition Informally, orientation of a cell complex K is preserved under a cell map g if g maps the oriented boundary of every non-collapsed cell oe 2 K into the oriented boundary of the image cell g(oe) 2 L. The sole purpose of this section is to express this statement precisely and in a computationally convenient form, using algebraic topological chains. Given a cell complex K; a p-dimensional chain, or simply p-chain, is a formal expression a an oe n ; where oe i are p-dimensional cells of K and a i are integer coefficients. Two p-chains on the same cell complex can be added together by collecting and adding coefficients on the same cells. The collection of all p-chains on K is a group denoted by C p (K) for 3. Using chains we can replace incidence, adjacency, and orientation computations with a simple algebra. In particular, we define the (algebraic) oriented boundary operation in terms of chains using only three coefficients from the set f\Gamma1; 0; +1g: (Such chains are often called elementary chains [1, 12].) Definition 7 (Boundary of a cell) The boundary of a p-cell oe is the (p \Gamma 1)-chain consisting of all (p \Gamma 1) cells that are faces of oe with coefficient +1 if the orientation of oe is consistent with the orientation of the face and \Gamma1 otherwise. 4 The coefficients are the simple algebraic means to compare the assumed orientation of a cell oe with the orientation induced in oe by an adjacent higher-dimensional cell (in the star of oe). For example, the oriented boundary of the 1-cell e 1 in Figure 6 is a 0-chain: @e 1 which implies that edge e 1 starts at vertex v 1 and ends at vertex v 2 . The boundary of an oriented 2-cell f in the same figure is a 1-chain: which tells us that the directions of edges e 1 and e 2 are consistent with the counter-clockwise direction of the face f; while the direction of the edge e 3 is not. The face f 2 in Figure 7 has a hole and its boundary is also a 1-chain: @f convention, the boundary of a 0-cell is defined to be 0. (a) Top view of K (b) Top view of L v3 v2 v8 v7 Figure 7: Top view of oriented boundary representation of the solid in Figure 2. The definition of boundary for an individual cell oe extends linearly to the boundary of any p-chain, i.e. an oe n is a p-chain, then it's boundary is a (p \Gamma 1)-chain with the usual rules of chain addition. In other words, the boundary operator is a linear function @ : 4 Here we rely on standard terminology in algebraic topology: 'face of oe' refers to any lower-dimensional cell incident on cell oe. The above definitions give precise characterization to the informal concept of "oriented boundary of every cell" as a chain. But the naming map g is a cell map and technically cannot be applied to chains; to see what g does to boundaries of cells, we need to extend g to maps on chains. Intuitively, a chain map g p takes individual p-cells (which are elementary p-chains) in K into individual p-cells in L - just as the cell map g does when the dimension p of the cells stays the same. But g may also collapse some p-cell oe 2 K onto a lower dimensional cell in L; since these lower dimensional cells do not belong to any p-chains on L, the chain map is instructed to simply ignore them by setting g p once again, we require that the action of a chain map on individual cells must extend linearly to arbitrary p-chains. 5 Together these three conditions give the usual definition of a chain map [17]: be a orientation preserving cell map. For each dimension p, define a chain map 1. if oe is a p-cell in K and g(oe) is a p-cell in L, g p 2. if 3. an oe n an Thus, any BR-deforming naming cell map g induces a family of chain maps g p in every dimension 2. It is also known [1, 5, 17] that every such family of induced chain map satisfies the following commutative diagram: \Gamma! C p (L) Since an oriented cell oe can be viewed as a chain with 0 coefficients attached to all other cells in the cell complex, the above commutative diagram can be enforced in a cell-by-cell fashion. Specifically, given an orientation preserving cell map g, the following simple condition is satisfied for each oriented p-cell oe 2 K: We will refer to this condition as the orientation condition, which is a precise restatement of the requirement that g maps the oriented boundary of every non-collapsed cell oe 2 K into the oriented boundary of the image cell g(oe) 2 L. 3.4 Examples revisited Consider once again the examples from section 1. Technically, every naming map g consists of the vertex, edge, and face maps. For brevity, we will only specify the action of g on those cells that are necessary for the presentation purposes. For example, below we assume that vertices are mapped to their implied images with a positive orientation and omit explicit description of the vertex map, even though the (chain of the) oriented boundary of every mapped edge is determined by the vertex map. Consider the naming map between the solids in Figure 2(a) and (c). The cells in the top view of the two solids are shown in Figure 7. The semantics of the update collapses the face f 1 and its boundary into the single edge e 0and its vertices. This deformation is reflected in the topological BR-deforming naming map g as follows: Notice that g is many-to-one and can be characterized as collapsing map. We already know from section 2.2 that g satisfies the star condition (2). Let's check if it also satisfies the necessary condition (3) in terms of induced chains maps g 1 and 5 In other words, chain map gp is a homomorphism [17]. A simple check shows that the necessary condition (3) @(g The rest of the cells remain topologically invariant under the naming map, and it is easy to check that the necessary condition (3) holds for all other cells in the boundary representation. Observe that the inverse map g \Gamma1 from the solid in Figure 2(c) to (a) is not BR-deforming; in fact it is not a valid cell map because it maps a one-dimensional edge e 0to a two-dimensional face f 1 . Nevertheless, such maps appear useful and are allowed in many systems. Therefore we will say that a one-to-many relation g is a expanding map if its inverse is a collapsing BR-deforming map. Now consider the apparent naming map between the solids in Figure 2(a) and (d): We already know that this map g is not BR-deforming (because it is not a cell map), and it is easy to verify that the necessary condition (3) does not hold either. The induced chain maps are: and the left and the right hand side of the equation (3) evaluate to different chains: Clearly, g is not collapsing map and the inverse map g \Gamma1 cannot be an expanding map. Next consider the example in Figure 3. We described two distinct naming cell maps g and h in section 2.2 that satisfy the star condition (2). We also argued that neither of them is appropriate - but for different reasons. We claimed that the naming map g computed by a commercial system is not BR-deforming because it is not orientation preserving. Indeed, g k (oe) = g(oe) for all cells in the boundary representation, because no cells have been collapsed, and in fact there is a one-to-one correspondence between the pre-edit and post-edit boundary representations. Yet, checking the necessary condition (3) for face f 3 , we see that Thus, @(g 2 same relation can be verified for all other faces in the boundary represen- tation; this shows that g is orientation reversing and is not BR-deforming. But now consider the necessary condition (3) for the second naming map h for the same solids: and so on. Notice that h reverses the assumed orientation of all edges connecting f 1 . Checking the necessary condition (3) on this oriented cell map: Similarly, one can check that map h does satisfy the necessary condition (3) for all other cells in the boundary representation and is a BR-deforming map associated with the 180 degree rotation of the solid about the X-axis. Furthermore, h is a cell homeomorphism and is geometrically BR-deforming because it preserves the combinatorial structure of the boundary representation, even though it may alter the geometry of some cells. Unfortunately, this particular BR-deformation is not consistent with the intended semantics of the parametric edit. In fact in this example, due to symmetry, there are many other geometric BR-deforming maps (e.g. rotation about Z-axis by 90 degrees). The guarantee of BR-deformation should not be confused with the guarantee of the correct semantics of the parametric update. Persistent naming 4.1 Sufficient conditions for an orientation preserving cell map In the typical scenario described in section 1, a pre-edit b-rep K and a post-edit b-rep L are given, and the system attempts to match the cells in K with the cells in L, i.e. to construct the naming map The matching process may or may not succeed. In terms of our definitions, there are at least three reasons why matching may fail: 1. A BR-deforming map g may not exist, because L does not belong to the same parametric family defined by K; 2. A BR-deforming map g does exist, but the system is not able to find it; 3. A BR-deforming map is not unique, and the system is not able to determine which map captures the semantics of the edit. 6 Whatever the outcome of the matching may be, the correctness of the result depends critically on the system's ability to determine if a constructed naming map is in fact BR-deforming. The formal machinery developed above is sufficient to determine if a given map g is an orientation preserving cell map, which is in turn a necessary condition for g to be BR-deforming. As our examples above have illustrated, these conditions are powerful enough to detect many (but not all) invalid edits in parametric solid modeling. To summarize, a map preserving cell map if and only if: 1. K and L are oriented consistently; 2. g is an oriented cell map; 3. if maps the oriented boundary of cell oe 2 K into the oriented boundary of the image cell g(oe) 2 L. The first requirement is enforced by assuming positive orientation for all vertices and faces (oriented counter-clockwise). The second requirement reduces to either the combinatorial closure condition (1) or the star condition (2). The third requirement is expressed by the orientation condition (3) which enforces the relative orientation between edges, vertices, and faces. All required conditions can be easily checked for a given map g; but it is important to understand the applicability and the limitations of the proposed conditions. Examples of Figures 3 and 6 may suggest that the orientation condition (3) is unnecessarily complicated because it detects only "global" orientation reversals as shown above. But imagine the solid in Figure 3 to be attached to a planar face of another solid (base), and now we can easily construct a local orientation reversing map which reflects a portion of the solid's boundary. Our experience suggests that such local reversals in orientation occur frequently in commercial systems, apparently due to partial and incremental updates in boundary representations. Enforcing the orientation condition for all cells would eliminate this type of invalid edit. The orientation condition (3) may appear to imply the star condition (2), since the closure of every cell includes its boundary points. To see why this is false, let us assume that g is an arbitrary (not necessarily cell) map that satisfies the first two conditions of the cell map (Definition 1). We can still use g to induce chain maps g p and use the boundary operator as before. Suppose we determine that the orientation condition (3) holds. Does this mean that g is an orientation preserving cell map? More specifically, does this mean that the star condition (2), or equivalently the combinatorial closure condition (1), is satisfied as well? Not necessarily. Consider the case when g collapses a p-cell oe 2 K to a lower dimensional cell g(oe) 2 L. By definition of a chain map, g p irrespectively of what g(oe) may be. Therefore, it is easy to construct a map g such that condition (3) is satisfied and yet g is not a cell map. Thus, it should be clear that the star condition and the orientation condition are not redundant for BR-deforming maps. On the other hand, suppose that dim(oe) for all cells in K. Since g preserves the dimension of every cell oe 2 K, the orientation condition (3) implies that 1-cells bounding every 2-cell oe 2 are in 1-to-1 6 In some cases, this is in fact the correct answer, because semantics of some edits has not been defined unambiguously. correspondence with 1-cells bounding g(oe 2 ), and 0-cells bounding every 1-cell oe 1 are in 1-to-1 correspondence with 0-cells bounding g(oe 1 ). In other words, g takes the closure of every cell in K to the closure of the corresponding cell in L, which implies the closure condition in the definition of cell map and the equivalent star condition (2). Furthermore, recall that when none of the cells are collapsed, the (geometrically) BR-deforming map g is a cell homeomorphism. It follows that the star condition (2) is redundant for geometrically BR-deforming maps. Thus, to check if a map is geometrically BR-deforming, it is only necessary that none of the cells are mapped to the cells of lower dimension and that the orientation condition (3) is satisfied for every cell. 4.2 Comparison with previously proposed naming methods Several heuristic methods have been proposed for constructing the naming maps g [10, 11, 18, 25, 31] in a variety of situations, including those where BR-deforming maps may not exist. A thorough analysis of all methods would not be practical in this paper, but it is instructive to check if the proposed techniques are consistent with our definitions when they do apply. The naming techniques proposed by Kripac[18] and Hoffmann et al [10, 11] rely on several common methods for naming cells in boundary representations. In particular, both start with the names of the primitive surfaces appearing in the parametric definition and construct the names for individual cells using: ffl the list of adjacent cells in the boundary representation; ffl the relative orientation of the cell with respect to the adjacent cells. In general, these two techniques are not sufficient to uniquely match cells in the two b-reps, and a number of additional techniques must be employed. For example, [18] also uses the history of editing to distinguish between cells, while [10] uses extended adjacency information and directional information associated with the particular solid construction method. The two naming techniques used by both [18] and [10] intuitively correspond respectively to the two types conditions that ensure an orientation preserving cell map. The use of adjacent cells is a means for extending the continuity of the naming map to the whole boundary of the solid, which relies on the combinatorial closure condition (1) and/or star condition (2). The matching of relative orientation of cells is basically equivalent to our orientation condition (3). On closer examination, it appears that [18] may map a cell oe 2 K to g(oe) 2 L when all cells adjacent to oe are mapped into cells adjacent to g(oe), without considering the orientation information. But as is clearly shown by example in Figure 3, such a naming map does not have to preserve orientation and may not be BR-deforming. By contrast, [11] matches oe to g(oe) only if the adjacent cells also preserve orientation. This method of naming will never result in orientation-reversing naming map. Both methods will find the geometrically BR-deforming map, if it exists and every cell has a unique name. It is not clear whether geomtric BR-deformation is always enforced when heuristic methods are used to resolve multiple matches between cells with the same names. Neither of the above methods are able to handle collapsing maps in general as defined in this paper, because both [18] and [11] require exact match of the adjacent cells, which in turn corresponds to the special case of exact equality in the star condition (2). For example, edge e 1 in Figure 2(a) is adjacent to faces f 1 and f 3 , while face f 1 is no longer present in the post-edit boundary representation of Figure 2(c). Thus, the proposed naming methods will never match e 1 to e 0, as they don't allow the strict inclusion relation implied by condition (2). More generally, all of the proposed methods assume that the dimension of the mapped cells will always be preserved, which is not the case for collapsing and expanding maps. Based on our results, it seems to make sense to redefine the adjacency conditions to correspond to the star condition (2). Rossignac observed that when every cell in a b-rep can be represented by a Boolean (set-theoretic) expression on primitive surfaces and halfspaces, the expression can be used as a persistent name for the cells [31]. In fact, it appears that some commercial systems use this method of naming in simple cases with limited success. 7 Boolean set-theoretic representations place no restrictions on the topology of the resulting 7 Constructing Boolean expressions for every cell in boundary representation is not trivial and may require a difficult construction known as separation [33]. boundary representations, imply a very different parametric family [34], and do not provide a mechanism for enforcing or even checking BR-deformation. A common method of matching cells in pre-edit and post-edit b-reps, which can be observed in some commercial systems, combines Boolean naming with the sequential ordering of the entities with respect to some external direction. For example, a system may attempt to distinguish between edges e 1 and e 2 in Figure 2(a) by indexing them along the length of the face (x-axis in the Figure). This method is most likely responsible for producing the solid in Figure 2(d) instead of (c): the identification of e 1 with e 3 changes the adjacency information so that e 0appears to be the "first" edge to fit the new name of e 1 5 Conclusion 5.1 Significance and limitations We have argued that the proposed principle of continuity and the implied notion of BR-deformation should be accepted as the basis for formally defining the semantics of a parametric family. Notice, however, that BR-deformation is not an equivalence relation: it is generally not true that boundary representation K can be deformed into L and vice versa. Thus, a parametric family of boundary representations can be defined in more than one way. For example, we could define a parametric family to include all those boundary representations that can be obtained by deforming one special 'master' boundary representation K. Alternatively, we could define two boundary representations to be in the same parametric family whenever one of them can be deformed into another. The latter definition may be more appealing because it does specify an equivalence class of solids, but deciding membership in such a class algorithmically appears to be nontrivial. The proposed formalism allows precise formulation and partial solution for a number of problems in parametric modeling. For example, the problem of "persistent naming" amounts to deciding whether two boundary representations K and L can be related by a BR-deforming map g that is consistent with the semantics of the parametric edit, and constructing such a map. A key computational utility for enforcing BR-deformation is the ability to decide if a given naming map between two b-reps K and L is BR-deforming. We have argued that any such g must satisfy the following two easily computable necessary conditions for every named (mapped) cell oe: ffl the combinatorial adjacency (star or closure) conditions expressed by (1) and (2); and ffl the orientation condition g p\Gamma1 Our formulations and the implied conditions do not completely solve the above problem, but they are the strongest possible, in the sense that the formulated problem could be solved only if solid modeling representations are enhanced with additional information that is not being used today. The main difficulty lies in the need to know what happens to the boundary representation not only for parameter values t 0 and , but also for all the values of t in the interval [t ]. For example, it is impossible to decide that the two boundary representations in Figure 5 do not belong to the same parametric range without considering the deformation process in E 3 . This is the very reason why we defined BR-deformation as a family of maps F t from E 3 into E 3 , even though the naming map takes jKj into jLj. To guarantee that a given naming map g is BR-deforming, we must show that g is an orientation preserving cell map and that it can be extended to a continuous map on the whole of E 3 . To be more precise, g will satisfy all the requirements of BR-deforming map if there exists a continuous orientation preserving map h from E 3 to E 3 , such that h restricted to jKj is a cell map g from jKj to jLj [1]. This sufficient condition for BR-deformation cannot be verified by considering only the structure of the two-dimensional boundary representations. Possible approaches to constructing BR-deforming maps include constructing the corresponding map h as a cell map on the three-dimensional decomposition of E 3 as suggested in [32, 34], or ensuring the properties of h by further restricting how g modifies individual cells [4, 36]. Other criteria for membership in a parametric family are possible. For example, one could require only that two boundary representations are related by an orientation preserving cell map. This would significantly enlarge the parametric family and would immediately establish computable sufficient conditions for persistent naming. These advantages would come at a considerable sacrifice: the loss of the physical principle of continuous deformation between solids in the same family. Nevertheless, this may be a reasonable and pragmatic compromise, given the difficulties of establishing BR-deformation as defined in section 2. 5.2 Semantics of more general edits It may appear that accepting the principle of continuity and BR-deformation as a basis for parametric modeling is too limiting because a number of practically important parametric edits cannot be described by BR-deformation. These include merging and splitting entities in boundary representations, (dis)appearance of holes, and common edits such as shown in Figure 8: as the slot moves to the left, the edge e 1 collapses into vertex v 1 . But this edit cannot be described by a cell map unless the edge bounded by vertices v 1 and v 3 is split into two edges, e 2 and e 3 , as shown in Figure 8(a). This would allow constructing a proper cell map that takes v 5 into v 0, v 4 into v 0, and so on. But of course, vertex v 2 would not be present in most boundary representations that rely on maximal faces [35]. Below we explain how our approach may be extended to formally define arbitrary parametric edits. (a) Original solid's b-rep K (b) Modified solid's b-rep L Figure 8: A parametric edit is defined by a splitting followed by a collapse While the notion of BR-deformation is not applicable to splitting and merging of faces and edges (such as needed in the example of Figure 8), recall that the main purpose of an orientation-preserving cell map was to enforce the continuity of the map in a cell-by-cell fashion; a natural extension of this principle would allow structural changes in the cells of K and L. In particular, subdividing or merging cells in a solid's boundary representation, e.g. as proposed in [23], has no effect on the boundary set itself (the underlying space). Thus, it seems reasonable that we can define a more general edit that preserves the spirit of continuous deformation by a sequence of splits, merges, and BR-deforming maps. Importantly, such edits include significant topological changes, including elimination of holes. Figure 9 illustrates a possible two-step procedure. First the collapsing map g takes the boundary of the hole c to a vertex c the resulting cell complex has a 2-cell that contains an isolated 0-cell in its interior. Then the subsequent merging step eliminates the 0-cell, without changing the underlying space. Finally, recall that the expanding map is not BR-deforming (only its inverse is), because it is one-to- many and is not a continuous cell map. Yet it appears that useful parametric edits require using such a collapse merge c c' Figure 9: Holes can be eliminated as a two-step procedure: collapse followed by a merge map. It may be feasible to define such edits by a sequence of maps such that each map in the sequence is collapsing, expanding, splitting, or merging. This approach may provide great flexibility in formally defining the semantics of parametric edits in a piecewise continuous fashion. For example, the reverse sequence of events shown in Figure 9 would allow the introduction of holes as a two-step procedure consisting of a split followed by an expanding map. This may also allow formally defining the semantics of feature attachment operations without using boolean operations. 5.3 Deformations of other representations The assumptions of continuity and continuous deformations make sense from an engineering point of view and can serve as a starting point for developing universal standards in parametric modeling. The proposed definition of BR-deformation and the implied necessary conditions can be employed in numerous other applications, including shape optimization, tolerancing, and constraint solving. For example, some heuristic methods proposed in [6] for identifying the correct solutions in constraint solving resemble the necessary conditions implied by our definitions, and the use of homotopy was also proposed in [19]. The principle of continuity is stated in terms of a particular representation scheme for solids. The same principle of continuity could be used to develop notions of parametric families with respect to other representation schemes. For example, it may be feasible to define a parametric family with respect to CSG representations [34], in which case the Boolean naming [31] may indeed correspond to a class of appropriately defined continuous deformations. The obvious and useful extension of our work is to general n-dimensional cellular models of pointsets, such as those represented by Selective Geometric Complexes [23] and more recently advocated by others [8, 26]. It should be intuitively clear that such cellular structures can be constructed and transformed into each other by combinations of expanding, collapsing, splitting, and merging maps. The cellular data structures (including boundary representations) are commonly supported by a number of so-called Euler operators [7, 20] that create, modify, and eliminate cells or collection of cells (loops, shells, rings, wedges, etc. Though the Euler operators are assumed to operate continuously on the underlying space of the cell complex, the continuity conditions are rarely enforced. The concepts of continuity, orientation preserving cell map, continuous deformation, and chain map apply to all cellular models without any modifications. Furthermore, the identified necessary conditions for deformation, including the star condition (2) and the orientation condition (3), hold as well. Acknowledgements This research is supported in part by the National Science Foundation CAREER award DMI-9502728 and grant DMI-9522806. The authors are grateful to Tom Peters, Malcolm Sabin, Neil Stewart, and Victor Zandy for reading the earlier drafts of this paper and suggesting numerous improvements. --R Algebraic Topology. Combinatorial Topology Polyhedral perturbations that preserve topological form. Basic Topology. A geometric constraint solver. Stepwise construction of polyhedra in geometric modeling. A geometric interface for solid modeling. Introductory Topology. Generic naming in generative On editability of feature-based design Dover Publications A road map to solid modeling. Morgan Kaufman On semantics of generative geometry representations. Parametric and Variational Design. Topology of Surfaces. A mechanism for persistently naming topological entities in history-based parametric solid models Solving geometric constraints by homotopy. An Introduction to Solid Modeling. Topology A First Course. Elements of Algebraic Topology. SGC: A dimension independent model for pointsets with internal structures and incomplete boundaries. Lectures in Algebraic Topology. Breps as displayable-selectable models in interactive design of families of geometric objects The generic geometric complex (GGC): a modeling scheme for families of decomposed pointsets. Mathematical models of rigid solid objects. Representations for rigid solids: Theory Representation of tolerances in solid modeling: Issues and alternative approaches. Issues in feature-based editing and interrogation of solid models CSG formulations for identifying and for trimming faces of CSG models. Maintenance of geometric representations through space decompositions. Separation for boundary to CSG conversion. What is a parametric family of solids? Alternative definitions of faces in boundary representations of solid objects. Sufficient condition for correct topological form in tolerance specification. --TR Introduction to Solid Modeling Geometric and solid modeling: an introduction Separation for boundary to CSG conversion What is a parametric family of solids? Solving geometric constraints by homotopy Polyhedral perturbations that preserve topological form The generic geometric complex (GGC) Representations for Rigid Solids: Theory, Methods, and Systems A Road Map To Solid Modeling --CTR Hilderick A. van der Meiden , Willem F. Bronsvoort, Solving topological constraints for declarative families of objects, Computer-Aided Design, v.39 n.8, p.652-662, August, 2007 Frank Y. Shih , Vijayalakshmi Gaddipati, Geometric modeling and representation based on sweep mathematical morphology, Information SciencesInformatics and Computer Science: An International Journal, v.171 n.1-3, p.213-231, 4 March 2005 Hilderick A. van der Meiden , Willem F. Bronsvoort, Tracking topological changes in feature models, Proceedings of the 2007 ACM symposium on Solid and physical modeling, June 04-06, 2007, Beijing, China Hilderick A. van der Meiden , Willem F. Bronsvoort, Solving topological constraints for declarative families of objects, Proceedings of the 2006 ACM symposium on Solid and physical modeling, June 06-08, 2006, Cardiff, Wales, United Kingdom Srinivas Raghothama , Vadim Shapiro, Models and representations for families of parts, From geometric modeling to shape modeling, Kluwer Academic Publishers, Norwell, MA, 2002 Thomas Convard , Patrick Bourdot, History based reactive objects for immersive CAD, Proceedings of the ninth ACM symposium on Solid modeling and applications, June 09-11, 2004, Genoa, Italy K. Y. Lee , C. G. Armstrong , M. A. Price , J. H. Lamont, A small feature suppression/unsuppression system for preparing B-rep models for analysis, Proceedings of the 2005 ACM symposium on Solid and physical modeling, p.113-124, June 13-15, 2005, Cambridge, Massachusetts David Marcheix , Guy Pierra, A survey of the persistent naming problem, Proceedings of the seventh ACM symposium on Solid modeling and applications, June 17-21, 2002, Saarbrcken, Germany Srinivas Raghothama, Constructive topological representations, Proceedings of the 2006 ACM symposium on Solid and physical modeling, June 06-08, 2006, Cardiff, Wales, United Kingdom Srinivas Raghothama , Vadim Shapiro, Consistent updates in dual representation systems, Proceedings of the fifth ACM symposium on Solid modeling and applications, p.65-75, June 08-11, 1999, Ann Arbor, Michigan, United States Guillaume Caumon , Charles H. Sword, Jr. , Jean-Laurent Mallet, Constrained modifications of non-manifold B-reps, Proceedings of the eighth ACM symposium on Solid modeling and applications, June 16-20, 2003, Seattle, Washington, USA Srinivas Raghothama , Vadim Shapiro, Topological framework for part families, Proceedings of the seventh ACM symposium on Solid modeling and applications, June 17-21, 2002, Saarbrcken, Germany

parametric editing;boundary deformation;cell complex;algebraic topology;persistent naming;soild modeling;boundary representation

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The edge-based design rule model revisited.

A model for integrated circuit design rules based on rectangle edge constraints has been proposed by Jeppson, Christensson, and Hedenstierna. This model appears to be the most rigorous proposed to date for the description of such edge-based design rules. However, in certain rare circumstances their model is unable to express the correct design rule when the constrained edges are not adjacent in the layout. We introduce a new notation, called an edge path, which allows us to extend their model to allow for constraints between edges separated by an arbitrary number of intervening edges. Using this notation we enumerate all edge paths that are required to correctly model the original design rule macros of the JCH model, and prove that these macros are sufficient to model the most common rules. We also show how this notation alows us to directly specify many kinds of conditional design rules that required ad hoc specification under the JCH model.

Introduction The technology used to manufacture integrated circuits [8] imposes certain limits on the sizes and relative positioning of features on the wafer. The resolution of optical lithography equipment, the amount of undercutting during a wet-etch processing step, the control over lateral diffusion and junction depth in dopant implant steps-all will set physical limits on how small device features can be, and how closely they can be separated, before the electrical behavior of the circuit changes. In addition, the interaction of successive processing steps-step-coverage problems in non-planar processes, or the ability to fill stacked vias, for example-also set limits on the manufacturability of the integrated circuits. Process and device engineers carefully characterize a semiconductor manufacturing process to understand these interactions, seeing them as complex statistical distributions of sizes and feature proximities and their effect on circuit reliability. Given only this raw data it would be very difficult and time-consuming to assess the overall reliability of a complete integrated circuit, or for a layout designer to draw a layout with a known level of reliability. Therefore, this statistical data is commonly codified into a set of much simpler design rules which can be given to the mask designers to guide them in producing a manufacturable layout [7]. Integrated circuit design rules abstract away the details of the process technology, and instead impose rules on the sizes and shapes of mask features and the separations and orientations of mask features in relation to one another. For example, a poly-separation rule might specify that the edges of unconnected rectangles on the polysilicon mask layer must be at least microns apart. One can also think of the design rules as imposing a set of constraints on the mask data-if all of the features in a given mask set conform to the design rules, the layout is considered to be design rule correct, and the designer can have some confidence that it can be manufactured with high yield. 1. Proof of design rule correctness is required before manufacturing can begin, so an important class of CAD tools, called design rule checkers, have been developed to check the mask data against the design rules in order to supply such a proof. In addition, layout generation tools, such as placement and routing tools, layout compactors, and leaf- less synthesis tools [3], require detailed information about the design rules so that correct layout can be produced. Therefore, in addition to making sure that the design rules are simple enough for mask designers to follow easily, it is important that they be specified in a formal and un-ambiguous manner so that they can be codified for use by algorithm designers. Two competing methods have evolved for the formal expression and modeling of VLSI design rules. The first employs a sequence of operations acting on the layout rectangles, such as the expansion and shrinkage of particular rectangle edges, and the union/intersection/inversion of overlapping rectangles. Design rules are defined by the presence of an empty or non-empty intersection between rectangles of a specified mask layer, or combination of mask layers, suitably expanded or shrunk by a given amount. Since these mask operations are global in nature, acting on all rectangles of a particular mask layer in the layout, we refer to these as mask-based design rule models. Mask-based modeling forms the basis of the commercial design rule checking software package Dracula [1], and has been formalized by Modarres and Lomax [5] through the use of set theoretic methods. We refer to the model of Modarres and Lomax as the ML model for design rules. The second method developed for the modeling of design rules treats the rules as constraints on the distance between individual edges of rectangles in the layout. There can potentially be a design rule constraint between any pair of non-perpendicular edges in the layout. The presence of a constraint, and if one exists the distance at which the edges are constrained to lie, depends on the layers making up the rectangles involved. Since this notion of design rules is local in extent, with its definition based on individual rectangle edges, we refer to these approaches as edge-based models. This second method has been formalized by K. Jeppson, S. Christensson, and N. Hedenstierna [2], using a modeling approach which we refer to as the JCH model for design rules. The two applications for design rules mentioned above, design rule checking and layout generation, expose two facets to the problem of design rule modeling. Mask-based methods seem dominant in the application of design rule checking, while edge-based models dominate in layout generation systems. In the former, the layout is static and unchanging, and we only seek to know if any design rule violations exist, and if so, where they are located. The global nature of the mask-based model suits this application well, and appears to lead the most naturally to elegant and efficient implementations based on well studied methods on computational geometry [6]. However, in the latter application one often requires the capability of making local queries in the layout as a new rectangle is added, to ensure that the rectangle is legally placed with respect to its neighbors. In addition, many layout generation applications, such as compaction, attempt to optimize the layout by finding optimal positions for each of the layout rectangles, subject to the design rule constraints. Because of its ability to represent the design rules as constraints between individual edges in the design, the edge-based model appears ideally suited to this application. In the course of our research into optimization based approaches to VLSI layout generation, we have studied the details of the JCH model, and have encountered several situations in which that model appears incomplete. In this paper we revisit the JCH model and provide some extensions which we feel contribute toward its universal application The basic notion of the edge-based design rule model, i.e. the definition of a design rule as a constraint on the relative positions of two non-perpendicular rectangle edges, is fairly simple. The complexity arises in attempting to characterize the specific rule, if any, which will apply to a given pair of edges. In the JCH model, each rectangle edge is assigned a type based on the layers which are present on either side. Given all of the different edge types, they exhaustively enumerate all pairs of edges which can appear adjacent to one another, and show how each possibility maps into a particular design rule. The strength of the JCH model is its ability to support this exhaustive analysis, which provides a proof of its correctness and completeness. This type of analysis has never been attempted for mask- based models, and it would probably be very difficult given the fact that all design rules are defined as global operations on the entire layout. In our study of the JCH model we have observed a shortcoming in the model as it was presented in [2]. The root of the problem is that, in their exhaustive enumeration of all possible design rules, the authors only examined the case of rules between adjacent edges in the layout. They show one example in which two edges, which should be constrained, are separated by an intermediate edge. Because of this intervening edge their model is unable to detect the proper constraint. Under several different interpretations of their definition for a design rule, they solve this problem by placing constraints on the values of the design rule parameters. When these constraints are satisfied the design rules that are missed will always be covered by other design rules, so there is no danger of missing a violation. However, we show one example in which their constraints are not strong enough, and we feel that constraining the model in this way is an inelegant solution to the problem. Instead, we feel it should be possible to correctly express these design rules so that they can always be detected. In this paper we extend the JCH model to cover, without ambiguity, situations under which design rule constraints can exist between edges with an arbitrary number of intermediate edges. To support the specification of design rules under this model we develop a simple grammar, called the edge-path grammar, which can be used to describe a more complete set of design rules. Using this grammar we describe the design rule macros of the JCH model and extend these macros to cover design rules which were un-representable in the notation used in [2]. 2. Previous Work-The JCH Model for Design Rules In the VLSI layout environment as defined by Mead and Conway [4], an integrated circuit is represented symbolically by a collection of geometric primitives on several different layout layers. The masks used to manufacture the circuit can later be derived from this collection of layout layers 1 . Formally, a layout layer represents a set of geometric primitives. All geometric primitives in a given set are said to be "on" that layout layer. One can also form derivative layers by performing various logical set operations on the layout layers. We will use the general term layer to refer either to a primary layout layer or a derivative layer. In the JCH model for design rules, the basic layout primitive is the rectangle. Each rectangle has four edges marking its boundaries. Where rectangles overlap, their edges are partitioned into disjoint intervals such that all points in a given interval have the same combination of layers on both sides. The rectangle partitioning operator defined by Modarres and Lomax [5] can be used for this purpose. Each edge interval (usually just called an edge) in a rectangle can be assigned a type based on the different layers that it separates. Furthermore, since each edge is normally part of only one layout rectangle (except in the case of "touch-edges", which we will address later), the two sides of the edge can be unambiguously defined as the inside layer and the outside layer. The type of an edge can therefore be expressed with the following notation. Edge-type: inside-layer/outside-layer (2-1) Figure 2-1 shows an example of two overlapping rectangles, one on layer A and one on layer B. The right-hand edge of the rectangle on layer A has been partitioned into three edge intervals with different edge types because of its overlap with the rectangle on layer B. To label the edges, the JCH model introduces the notation shown in the figure. We consider A to be a boolean variable indicating whether one side of an edge is on layer A or not. If A appears uncomplemented it indicates that side is on layer A. Conversely, if A appears complemented (written ) it indicates that side is not on layer A. Therefore the edge-interval labeled , marked with a bold line, indicates that its inside-layer consists of both layer A and layer B, while its outside-layer consists only of layer B. 1. In the interest of reducing the number of layout layers, it may not be necessary to maintain a direct one-to-one correspondence between the layout layers and the masks. For instance, it is common to use only a single layout layer to represent interconnect vias, as the specific oxide-etch mask that a via will appear on can be inferred from the surrounding layers. layer A AB/AB edges AB/AB edges AB/AB edges edge-interval Figure Two overlapping rectangles and the resulting edge partitions and types. AB/AB edges A Edges and edge intervals are represented geometrically by a line in the plane. In manhattan layouts, edges are either vertical or horizontal. Vertical edges have a single x-axis coordinate, labelled , and a pair of y-axis coor- dinates, labelled and , where by convention . Similarly, horizontal edges have a single y coordinate and a pair of x coordinates. For simplicity, all of our examples will assume vertical edge, but they are trivially extended to horizontal edges as well. In the JCH model, there are two classes of design rules, both involving rectangle edges. Forbidden-edge rules simply forbid a particular edge-type from occurring. Restricted-separation rules specify a separation value to be enforced between two parallel edges 2 . Restricted-separation rules can be viewed as a constraint enforcing either a minimum, maximum, or exact separation between two edges, edge1 and edge2, and can be written as follows in the case of vertical edges. Here we are specifying that the distance between the two edges, in the direction perpendicular to the edges, is constrainted by some value d. The variable d is the design rule parameter, and is in general a conditional quantity whose value may depend on any properties of the two edges that can be extracted from the layout. Common examples of these properties are: . the edge-types of edge1 and edge2 . the presence of rectangles overlapping edge1 or edge2 . the lengths of edge1 and edge2 . the widths of the rectangles formed by edge1 or edge2 . the edge-types of edges that lie between edge1 and edge2 . the degree of overlap of the two edges in the direction perpendicular to the edges This could imply a semantically complex set of rules, but most common cases tend to be simple. The JCH model advocates the creation of macros to express the common cases, and the treatment of more complex cases as further conditions that can be placed on the design rules when required. The JCH macros involve only the properties in the first bullet, the edge types of the two edges involved. Rectangles overlapping one or both of the edges can conditionally form a special class of rules called "conjunctive" design rules, and properties of the edges such as their length form classes of so-called "conditional" design rules in the JCH model. We will review these in Section 2.3. In this paper, we will show several reasons to include the property of the fifth bullet, the edges types of all edges lying between the two edges in the layout. The property mentioned in the last bullet is a common implementation detail worth mentioning. Often, it is only necessary to enforce a design rule between two edges that overlap somewhere along their length (i.e. vertical edges must overlap in the y direction) as shown in Figure 2-2(a). However, in some design rules (typically the spacing rules) the notion of overlap must be implemented carefully in order to ensure that the rule is enforced correctly at rectangle corners. In these cases, for the purpose of detecting the edge overlap, the ends of the edges must be extended by the design rule distance d, as shown in Figure 2-2(b). In the JCH model this edge extension is controlled by the "con- cave" modifier to their "edge function". At these extended rectangle corners, some foundry design rules require the design rule constraints to enforce the manhattan distance between the edges, and in some cases they allow the shorter euclidean distance, as shown In Figure 2-2(c). The latter case can be supported by making the design rule parameter d conditional on the y coordinates of vertical edges. The parameter d can be reduced as the edges cease to overlap in the y direction, and the constraint is eliminated altogether when the edge-extensions cease to overlap. In order to describe a design rule, the value of the design rule parameter must be given along with the conditions under which it applies. As we mentioned above, in the JCH model the value of the design rule parameter in restricted- 2. Here we are assuming a manhattan model for layout in which all edges must be parallel to one euclidean axis. The model can easily be extended to non-manhattan layouts by allowing constraints between all edges that are not perpendicular. edge x edge y1 edge y2 edge y1 edge y2 separation rule - maximum separation rule exact separation rule separation constraints is commonly conditional only on the edge-types of the pair of edges involved. This condition is given a notational representation which we refer to as an edge-pair, written in the following way: Since the two edges in the edge-pair notation share a common outside layer, this implies that constraints can only be placed between adjacent edges. At the end of Section 2. we will show that restricting the edges to be adjacent proves to be a serious limitation. We will expand the definition of the design rule to include not only the edge-types of both edges, but also the types of edges that may appear between these two edges. For brevity in our expanded defini- tion, and in order to emphasize that the two edges must share a common outside-layer, we introduce the following notation to represent the same edge-pair shown in (2-3). We call this expression an edge-path, as it represents a sequence of edges between which the constraint lies. We will develop the grammar for edge-paths more fully in Section 3. In the remainder of this section we will describe the design rules of the JCH model using the edge-path notation. These rules can be split up into two categories, one which contains rules involving only a single layer, and one which contains rules involving two layers. The former are referred to as intra-layer rules, the latter inter-layer rules. These will be reviewed in Sections 2.1. and 2.2. Section 2.3. will discuss the author's method for addressing design rules which involve more than two layers, and Section 2.4. will address the limitations of the JCH model. 2.1. Intra-Layer Design Rules Some design rules involve only a single layer. These are referred to as the intra-layer rules. When checking intra- layers rules, one layer is considered at a time and the presence of all other layers is ignored. If we refer to this layer as layer A, the layer type on either side of an edge can take on only the values A and . There are thus four possible edges that can occur in this situation: Of these, the patterns and do not represent actual edges, as the layer is the same on both sides. Therefore d constraint euclidean distance manhattan distance d Figure Conditions placed on the design rule constraint by the degree of vertical overlap of two vertical edges. In (a) a constraint is required because the rectangles overlap in the y direction. In (b) a constraint is still required to prevent design rule violation at the rectangle corners until the rectangles are separated in y by more than the distance d. In (c) is demonstrated two possible interpretations of design rule constraints at rectangle corners. (b) (c) d d overlap (a) A A A A A A A only the remaining two edges will appear in the design rule constraints. Since the design rule constraints must share a common outside-layer, there are only two intra-layer constraints: In (2-6) each row represents one design rule. The expressions WIDTH and SPACING represent two different design rule parameters which correspond to the parameter d in (2-2), and the associated edge-path indicates the conditions under which this design rule is applied. These two intra-layer rules are diagrammed in Figure 2-3. 2.2. Inter-Layer Design Rules The case when two different layers can be involved in a design rule is more complex than the single-layer case. In the JCH model these are referred to as the inter-layer design rules. For any pair of layers present in the design, if we refer to these layers as layer A and layer B, the layer types on either side of an edge can take on the values , , , or . There are therefore 16 possible edges that can be formed, as shown below. As in the intra-layer case, edges with the same layer type on both sides are not considered to be true edges, and we can therefore eliminate four of these "non-edges" from the list shown in (2-7). In addition, there are four edges in which both layers undergo a transition, which can also be eliminated. These edges are called "touch edges" in the JCH model, and they will not normally be encountered during inter-layer design rule checks. We elaborate on this issue in Section 4.3. Only the eight edge types shown in bold in (2-7) will be retained for the following analysis. There are four possible layer types that can represent type1 in a design rule constraint. In order to form a valid pair of edges there are only two possible values for the type2 layer given type1, and only two possible values for type3 given a value for type2. This results in 16 different inter-layer constraints, as shown below. A A A A layer A layer A layer A WIDTH SPACING Figure 2-3: The two intra-layer design rule parameters of the JCH model In (2-8) notice that in the SPACING and WIDTH rules the variables for either layer A or layer B appear with the same inversion in all three of the type1, type2 and type3 layer fields. Thus, this layer is either absent on both sides of both edges, or always present on both sides. These edge paths will also be matched by the intra-layer WIDTH and SPACING rules, and the JCH model assumes that they therefore need not be checked here. With the WIDTH and SPACING rules eliminated, only the remaining four rules require inter-layer design rule parameters. These four rules are illustrated in Figure 2-4. Note that in the JCH model the A-EXTENSION-OF-B rule is given the name "extension" while the B-EXTENSION-OF-A rule is given the name "margin". We feel that these names are ambiguous and have chosen more descriptive names. Here we would like to point out that neglecting the SPACING and WIDTH constraints in the inter-layer macros makes sense in the case of the first column in which the unchanging layer is always absent. However the case of the second column, in which the unchanging layer is present and overlaps both edges, is an example of a "conjunctive" rule, as discussed in Section 2.3. Conjunctive SPACING and WIDTH rules may be different in some processes than the non-conjunctive rules, in which case a user may wish to retain these edge-paths as legitimate design rules. 2.3. Conditional and Conjunctive Design Rules Two classes of design rules which cannot be directly represented in the JCH model are examined by its authors. They refer to these two classes as conditional and conjunctive design rules. Conditional design rules are ordinary intra-layer or inter-layer design rules in which the constraint parameter is conditional on some outside factor. An example of this is the common metal-halation rule which requires a larger separation between metal wires when one of the wires is wider than some threshold. This rule is a result of processing effects due to the high reflectivity of interconnect metal and its effect on photoresist exposure times. The specification of this rule requires that the SPACING rule for a particular layer be made conditional on the width of the rectangle to which one or both of the edges belongs. This width would have to be extracted from the layout prior to design rule constraint generation. In conjunctive design rules, the constraint parameter value depends on the presence or absence of an otherwise unrelated layer. As an example, step coverage problems in non-planar metallization processes may require the SPACING between two metal-2 wires to be increased if either wire overlaps metal1. This is an example of the conjunctive CLEARANCE A-EXTENSION-OF-B B-EXTENSION-OF-A OVERLAP Figure 2-4: The four inter-layer design rule parameters of the JCH model inter-layer SPACING rule discussed in the previous section. More complex conjunctive design rules may involve a third layer, as in a conjunctive CLEARANCE rule between metal-1 and metal-2 wires which is increased if either wire overlaps a polysilicon wire. This latter case, in which an inter-layer design rule between two layers A and B is conditional on a third layer C is the only situation in which a design rule may involve edges between more than two layers. The authors suggest that some conditional and conjunctive design rule checks can be accomplished by forming carefully chosen derivative layers. When this fails they suggest an ad-hoc method in which the design rules are extended with special conditional constraints tailored to the specific problem. Conditional design rules are difficult to characterize, as there can be an almost endless variety, and it may be the case that one will always have to resort to ad-hoc methods. However, the extensions to the JCH model that we propose in Section 3. are capable of expressing any conceivable conjunctive design rule. 2.4. Limitations of the JCH model Recall that, in the JCH model, a design rule is represented by a constraint between a type1/type2 edge and a type3/ edge. This constraint can be interpreted and implemented in a number of ways. The authors choose to interpret the constraint as a region of width d extending from the type1/type2 edge from which the type3/type2 edge is forbidden from appearing, and vice versa. For reasons of efficiency, the authors have chosen to implement their system with a simplified version of this inter- pretation. Instead of searching for the type3/type2 edge within the constraint region, they search for any instances of the layer type3. This would initially appear to be equivalent, but they discuss a case, shown in Figure 2-5(a), in which these interpretations are not consistent. In this case, their simplified model would search for the layer inside the constraint region of the edge, and would identify the clearance constraint shown. This should not be flagged as an error since it does not comply with the original definition of clearance. They state that, in the case of the clearance constraint, their simplified interpretation is only consistent with their original definition if the following constraint holds. A A spacing extension overlap clearance "correct" clearance A A spacing extension overlap clearance "correct" clearance (a) (b) Figure 2-5: Examples of a clearance rule as identified by two different constraint implementa- tions. Figure (a) shows the result of searching for the type3 layer, layer , inside the constraint region of the type1/type2 edge. Figure (b) shows the result of searching only for the layer that undergoes a transition across the rule definition's type3/type2 edge, in this case layer . In both figures we show the "correct" clearance rule that we identify by eye. { When saying that the simplified interpretation is "consistent" under these conditions, they mean that edges identified as constraints by the simplified interpretation, but not the original design rule definition, will never be flagged as errors unless they violate other design rules as well. Similar constraints can be derived for the other three inter-layer parameters. However, we would like to point out that under their simplified interpretation, what most designers would consider the correct clearance rule, marked in Figure 2-5(a), is never checked. This could lead to a design rule check passing when in fact an error exists. Furthermore, and potentially more serious, the correct clearance rule is not identified by the authors original definition of design rules either. The reason for this failure is the presence of an intermediate edge between the edges that should be checked, which causes the outside layers of the two edges to be different. The following constraint, more stringent than that of (2-9), is required to guarantee that clearance violations are not missed in this example. Again, similar constraints are required for the other three inter-layer parameters. Interestingly, the authors present a second simplified interpretation of the design rule constraints which actually detects the correct edges in this case. An example of this interpretation in the context of the same clearance rule used in Figure 2-5(a) is shown in Figure 2-5(b). Instead of searching the constraint region of the type1/type2 edge for instances of the layer type3, they search only for the layer that undergoes a transition across the type3/type2 edge (recall that, in the absence of touch edges, only one layer will transition across the edge.) In this example it is equivalent to searching for instances of layer B inside the constraint region of the edge. In this situation, the check does in fact identify the correct clearance constraint, as shown. The authors indicate that this second simplified interpretation for the design rules is only consistent with their original definition for design rules given the same constraint written in (2-10). What they failed to observe is that when this constraint is not met, it is the simplified interpretation that is correct and not the original interpretation. The adoption of this second simplified interpretation may seem to solve the inaccuracy problem, but in Section 3.3. (Figure 3-4) we will show an example of a clearance rule that would not be located by this method either. As a final comment on this second simplified interpretation of the design rule constraints, we note that, unlike the check in Figure 2-5(a), the check in Figure 2-5(b) is not reflexive. It is not possible to identify the same constraint by searching for layer A from the edge, as the only inter-layer constraint that begin on edges is the A-EXTENSION-OF-B constraint. Until this point, all constraints have had duals which allowed the constraint to be identified starting from either edge involved. When all rules are reflexive, design rule checking can be simplified because the layout only has to be scanned in one direction along each manhattan axis. In the following section we develop our edge-path grammar and show how it provides an unambiguous definition for the design rules under all circumstances. We also show how the edge-path grammar can be used to extend the JCH design rule model to include a straightforward and consistent model for conjunctive design rules. 3. Extending the JCH model The true nature of the limitations which have been demonstrated in the JCH model is its inability to model situations in which one or more intermediate edges lie between the two which must be constrained. When this occurs, the two edges no longer share a common outside layer, and the rule cannot be written as in (2-3). What is required is a more general view of the design rule constraints that spans multiple edges. In order to express these, we will extend the notation which we presented in (2-4) of Section 2., which we have called the edge-path grammar. 3.1. The Edge-Path Grammar When defined in the most general way possible, a design rule constraint can exist between any two non-perpendicular edges in a circuit layout. There can be an arbitrary number of intermediate edges between the two constrained edges, and the value of the constraint can depend on the exact pattern of layer boundaries that make up these edges. If { we choose a point on each of the pair of edges we wish to check, and connect them with a line as shown in Figure 3- 1, this line will originate at one edge, cross a specific sequence of layers as we move from one edge to the other, and terminate on the second edge. We call this sequence an edge-path, which is written as follows. In an edge-path, type1, type2, ., and typen represent boolean products of layer variables which indicate the presence or absence of each layer in that particular region. We will now present this in a more formal way. Definition 3-1: l is a layer in an integrated circuit layout. The layer l can represent either a primary layout layer, or a derivative layer formed by a sequence of logical set operations on the layout layers. Definition 3-2: L is the set of all layers present in a particular integrated circuit layout. Definition 3-3: p is a boolean variable representing a layer . The range of p is the set , in which the value indicates the absence of layer l, and the value indicates the presence of layer l, in a particular region of the layout. Alternatively, the range of p can be written where and . If a variable for a particular layer does not appear in an expression, it can be considered a don't-care condition. 3-4: P is an expression representing the conjunction of one or more boolean variables . P represents the combination of layers from the subset which are present in a particular region of a layout. Since the variables p have a binary range, P can be thought of as a cube in an n-dimensional boolean lattice. Definition 3-5: An edge-path E is a totally-ordered set , written , which represents a sequence of layers encountered along a line drawn in the plane of the layout. Each pair of adjacent members in the set, for all , represents an edge at which one or more layers undergoes a transition. Definition 3-6: The length of an edge-path E is defined as , the number of edges encountered when following a particular edge-path. Using the edge-path grammar we have a more powerful notation for representing VLSI design rules than that of the JCH model's edge-pair. We have already demonstrated that we can represent all of the JCH model intra-layer and inter-layer rules as edge paths of length two. However, we have also demonstrated that this set of edge-paths is inadequate to describe the design rules under all circumstances. We will show that a complete set of design rules, including conjunctive rules and rules involving touch edges, can be modeled under the edge-path notation if the edge paths are allowed to become arbitrarily long. Fundamentally, the design rules for a semiconductor process can be described as a large set of edge-path/parame- ter pairs. However, in order to simplify design rule entry, it is advantageous to follow the methodology of the JCH model and classify these edge-paths into sets which effectively represent the same design rule. This facilitates the use of design rule macros that free the user from an intimate knowledge of the full model. Many such classifications are possible. We will demonstrate that any edge-path can either be classified as one of the six parameters of the JCH model (SPACING, WIDTH, CLEARANCE, A-EXTENSION-OF-B, B-EXTENSION-OF-A, and OVERLAP), or it corresponds to a configuration of geometry that does not normally require a design rule. In some situations this simple classification will be inadequate. It neglects the conjunctive design rules and rules involving touch edges. In some cases it may actually place a design rule on a geometrical pattern that should be forbidden from occurring. Obviously the edge-path classification that we present should only be considered a typical default. The user should be allowed to override the macros and forbid particular configurations, or add new edge- paths for situations that are not covered, and we will discuss several situations in which it will be common to do so. Figure 3-1: A edge-path between two edges with an arbitrary number of intervening edges l L { }{ } 1 { } { { } { } { } L { } 3.2. Edge-Paths of Length Three We begin by studying all edge-paths of length three. For simplicity, we have decided to forbid touch-edges for the same reasons as stated earlier, so there are 32 edge paths of length three. It is easy to show that all non-conjunctive intra-layer design rules can be captured in sequences of length two, so we will only concern ourselves with the inter-layer design rules. Recall that our methodology is to examine all possible edge-paths and show that each one either matches a specific JCH model parameter, or it corresponds to a situation that is not a traditional design rule. Half of the 32 edge-paths of length three fall under this latter case and do not match any of the six JCH model parameters. These are shown below, and several examples are diagrammed in Figure 3-2. One thing that each of the 16 edge-paths in (3-2) have in common is that, like the examples in Figure 3-2, they begin on an edge of a rectangle, cross the near-side of a second rectangle, and terminate on the far edge of the second rectangle, or vice versa. These two edges are usually constrainted by the sum of a width and a separation/clearance/ extension rule, and we know of no design rules in use that specify a more stringent rule in these circumstances. For this reason, we choose to ignore these patterns when they are matched in a layout, essentially making them non-con- straints. However, we emphasize that a designer could override this decision and specify design rule parameters for each of these patterns if desired. Each of the remaining 16 edge-paths can be matched to a specific intra-layer or inter-layer design rule parameter as defined in the JCH model. These are shown below and diagrammed in Figure 3-3. A A Figure 3-2: Examples of edge-paths of length three that do not correspond to any inter-layer rule AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB AB Upon examining the 16 edge-paths in (3-3), we note that, just as in the case of the inter-layer edge-paths of length two, the SPACING and WIDTH rules would be matched by the corresponding intra-layer design rules, and we can eliminate them from the default inter-layer design rule specification. However, as before, these patterns represent conjunctive design rules and the user may wish to specify these edge-paths to the system manually. The remaining eight edge-paths should be generated by the inter-layer design rule macro in addition to the edge paths in (2-8). They represent layout patterns which may exist in the layout but which are not matched by any of the design rules in the JCH model. This will eliminate the constraints under which the JCH model will either indicate a false positive or false negative during a design rule check as a result of layout situations like those addressed in [2], and repeated in Figure 2-5. Of course some of these patterns, especially the new EXTENSION edge-paths, may correspond to geometry which should not occur in the layout. If this is the case, the designer should specify these edge- paths as forbidden edge types instead of allowing them to be matched to the corresponding inter-layer design rule. 3.3. Edge Paths of Length Four and Above There are 64 edge paths of length four, 128 edge paths of length five, with the number continuing to double as the length increases. For brevity we will only summarize our results in this section. It turns out that we can identify conjunctive WIDTH and SPACING constraints with edge-paths of arbitrary length. We show the four layer A WIDTH and SPACING design rules of length four in Figure 3-4. It is clear from these examples that we can insert an arbitrary number of layer B edges between the layer A edges whose width or spacing is to be constrained, and vice versa. How- ever, all of these cases can probably be covered by a single conjunctive intra-layer macro, if not by the original non- conjunctive intra-layer macro. Also shown in Figure 3-4 are the only inter-layer design rules of length four or larger, two CLEARANCE rules. These two edge-paths, shown below in (3-4), should also be generated by the design rule compiler when an inter-layer design rule macro is specified by the user, or specified as forbidden edge-paths by the user if they correspond to an illegal configuration. It is interesting to note that the CLEARANCE rules shown here would not be detected by the JCH model's second simplified interpretation for the design rule constraints, which was discussed in Section 2.4. Under this interpretation the CLEARANCE rule is detected by searching for instances of layer A within the constraint region of edges and instances of layer B within the constraint regions of edges. However, there are no regions along this edge path. A A A A-WIDTH A B-WIDTH A A A A Figure 3-3: The 8 inter-layer design rules for edge-paths of length three A A Just as was done in Equation 2-9, we can write the following constraint on the design rule parameters under which the original JCH model will still be valid for these two CLEARANCE rules. This constraint is more strict than the constraints described by the authors, and it corresponds to a situation in which their model could potentially have generated false positive design rule checking results. As a final result, we state that, with the exception of the conjunctive intra-layer design rules mentioned above, which continue to appear in edge-paths of any length, there are no inter-layer design rules of length greater than four. It should be obvious that OVERLAP rules cannot exist with path lengths greater than two (i.e. there can be no intermediate edges.) In the case of the CLEARANCE and EXTENSION rules, except in the cases we have already dem- onstrated, long edge-paths always contain shorter sub-strings that also match the same rule. This shorter sub-path therefore covers the longer edge-path, guaranteeing that it will never be in violation. Examples of this effect are presented in Figure 3-5. A A A-WIDTH A A A A A Figure 3-4: The SPACING and WIDTH design rules for layer A (the rules for B are the same but with A and B reversed), plus the two CLEARANCE rules, for edge-paths of length four. A A A A-WIDTH { A A Figure 3-5: Example of cases in which a design rule with a path-length of four is blocked by an identical rule which is a sub-expression of the same edge-path A A 4. Implementation Suggestions In this section we discuss several implementation details which are important if this model is to be incorporated into a practical design system. In Section 4.1. we present a simple state machine representation of the design rule edge paths which neatly summarizes the results of Section 3. As mentioned previously, there are no conjunctive rules in the default specification, and by default touch edges are forbidden between layers that are to be checked for viola- tions. In Sections 4.2. and 4.3. we show how this system can be extended to specify conjunctive inter-layer design rules, and how touch-edges can be handled in an elegant way. 4.1. State Machine Representation of Design Rules Section 3. presented a large collection of edge-paths and the design rule parameter that is associated with each. These can be neatly summarized by the state diagram shown in Figure 4-1. We can view an edge-path as a straight-line drawn in the layout plane, beginning at a point on a rectangle edge (the source edge), and projecting in a direction perpendicular to that edge. As this line is traversed from the source-edge outward, each new edge that is encountered will cause a state transition that maintains a record of the previous edges that have been encountered, and indicates when a design rule check is needed. The line traversal can end when a terminal vertex in this state diagram is reached. Design rule constraints are required when certain states are reached, as shown in the state diagram. This conceptual view, in addition to providing a compact summary of the edge-path checks, leads to a direct implementation of the common shadow algorithm used for constraint-generation in one-dimensional compaction, a good summary of which is presented in [9]. begin BEA BEA C O BEA BEA O done done done done done done done done done done done done done done Figure 4-1: State machine that summarizes the edge-paths resulting in design rule checks. Key: 4.2. Conjunctive Design Rule Specification We have already demonstrated how four of the inter-layer edge-paths of length two correspond to rules that can be considered conjunctive versions of the intra-layer width and spacing rules. These edge-paths are ignored in the JCH model because they are covered by the intra-layer rules. However, if conjunctive versions of these rules are required, some method must be present to allow them to be checked. Conjunctive versions of the inter-layer rules are also common. These can be specified simply by allowing edge- path expressions to contain three or more boolean layer variables. As an example, Figure 4-2 shows a conjunctive CLEARANCE rule in which the separation between layers A and B changes because of their overlap of layer C. If conjunctive rules are required, the JCH model outlines a strategy which makes use of derived-layers obtained through logical set operations on the existing layout layers. However, in the design rule system that we have outlined, these rules can be specified directly by adding their edge-paths as additional checks. Of course, when three or more layers are involved in the edge path expressions, this will result in a large number of possible edge-paths that a user would be required to enter into the system, and that must be added to the state machine in Figure 4-1. 4.3. Touch Edges In the JCH model, "touch-edges" are edges at which more than one layer undergoes a transition. Figure 4-3 shows the two types of touch-edges which can occur during inter-layer checks. Outside-touch edges ( and are formed when rectangles touch but don't overlap, and inside-touch edges ( and are formed when rectangles overlap and touch on one edge. The authors of the JCH model chose not to include touch edges in their definitions of the inter-layer rules. Instead they introduce two flags into the inter-layer macro that, if set, simply forbid them from appearing for each pair of lay- ers. Their reasoning is not given, but straightforward. It is usually the case that if any design rules exist between two layers, touch edges will be forbidden from occurring. If the two layers interact during processing so as to require design rules between them, the overlap, extension, and clearance rules ensure that rectangle edges are separated. It is inconsistent under these circumstances to allow two rectangles to touch without overlapping. Conversely, if there are no design rules between two layers, implying that they don't interact with each other during processing, there is no reason for touch edges to be forbidden. Figure 4-2: A conjunctive CLEARANCE rule involving layers A, B, and C. A A OUTSIDE_TOUCH INSIDE_TOUCH Figure 4-3: Examples of OUTSIDE_TOUCH and INSIDE_TOUCH edges It is unclear exactly how the design rules should be defined when touch edges are allowed on layers between which design rule are defined. One valid option is that the designer could enter edge-path expressions to specify the design rules between touch-edges and other edges. However, this could result in a large number of new rules. The JCH model deals with touch edges by considering the inter-layer rules to be "minimum distance or touch rules", though they fail to elaborate precisely what this means. We interpret this as meaning that when a touch-edge is encountered, constraints are generated to the touch-edge exactly as if it were two separate edges of each layer individ- ually. Thus, there can potentially be two constraints generated between a normal edge and a touch-edge, and there may be up to four constraints generated between two touch edges. If we make use of the state-machine description of the design rules shown in Figure 4-1, there is a particularly elegant way of viewing constraint generation for touch edges under this "minimum distance or touch rule" interpreta- tion. If a touch edge is encountered as an edge-path is being traced through the layout, the state machine simply explores both branches out of the current state. 5. Conclusions In Section 2.4. we discussed several limitations of the JCH model in the context of several examples provided by its authors. We showed how the original definition of the design rules, a constraint between a type1/type2 and a type3/ type2 edge, can fail to match some edge patterns to the correct design rule. Two alternative interpretations for the design rule definitions are presented by the authors. The first, which they chose to implement, searches for any regions of type3 material inside the constraint region of the type1/type2 edge, and vice versa. The second interpretation presented by the authors searches the type1/type2 constraint region for the presence of the layer in type3 that makes a transition across the type3/type2 edge. As the authors point out, all three interpretations can yield different results under certain conditions. All three methods can report false design rule violations under certain conditions given by the authors, and more importantly, it can fail to report design rule violations under certain conditions, as we have shown. The authors of the JCH model have presented some constraints under which their various implementations of the design rules are correct, in the sense that a design rule violation will not be reported between edges which are incorrectly flagged with constraints by their model. However, we have demonstrated that their constraints are insufficient, causing legitimate design rules to be missed. With tighter constraints their model can easily be corrected, but we do not think that creating conditions under which the model does not apply is an elegant solution. The root of the limitations of the JCH model are its inability to recognize design rules between rectangle edges which are not adjacent. We propose an extension to the JCH model in the form of a new syntax for expressing the lay-out patterns that define the different design rules, which we call an edge-path. Edge-paths can express a constraint between two edges that are separated by a sequence of edges of arbitrary length. With this syntax, we can perform an exhaustive examination of all possible edge-paths and check each pattern to see if it corresponds to a design rule that should be enforced. We have demonstrated that we can characterize all edge paths through the use of the six original JCH model design rule parameters. The two intra-layer parameters can be expressed with only four edge-paths, all of length two. The four inter-layer parameters can be expressed using edge paths with a length of at most four. We have also shown how this new syntax can be used to elegantly express design rules involving "touch-edges" as well as the conjunctive design rules mentioned by the authors of the JCH model. Acknowledgments The authors would like to thank Ronald Lomax and Jeff Bell for reading early drafts of this paper and for their insightful comments. --R Cadence Design Systems Formal definitions of edge-based geometric design rules Combinatorial Algorithms for Integrated Circuit Layout. Introduction to VLSI Systems. A Formal Approach to Design-Rule Checking Computational Geometry: An Introduction. A statistical design rule developer. Silicon Processing Symbolic layout and compaction. --TR Computational geometry: an introduction Combinatorial algorithms for integrated circuit layout Introduction to VLSI Systems Magic Lyra

design rules;layout verification;design rule checking

293906

Discrete Lotsizing and Scheduling by Batch Sequencing.

The discrete lotsizing and scheduling problem for one machine with sequence-dependent setup times and setup costs is solved as a single machine scheduling problem, which we term the batch sequencing problem. The relationship between the lotsizing problem and the batch sequencing problem is analyzed. The batch sequencing problem is solved with a branch & bound algorithm which is accelerated by bounding and dominance rules. The algorithm is compared with recently published procedures for solving variants of the DLSP and is found to be more efficient if the number of items is not large.

Introduction In certain manufacturing systems a significant amount of setup is required to change production from one type of products to another, such as in the scheduling of production lines or in chemical engineering. Productivity can then be increased by batching in order to avoid setups. However, demand for different products arises at different points in time within the planning horizon. To satisfy dynamic demand, either large inventories must be kept if production is run with large batches or frequent setups are required if inventory levels are kept low. Significant setup times, which consume scarce production capacity, tend to further complicate the scheduling problem. The discrete lotsizing and scheduling problem (DLSP) is a well-known model for this situation. In the DLSP, demand for each item is dynamic and back-logging is not allowed. Prior to each production run a setup is required. Setup costs and setup times depend on either the next item only (sequence independent), or on the sequence of items (sequence dependent). Production has to meet the present or future demand, and the latter case also incurs holding costs. The planning horizon is divided into a finite number of (short) periods. In each period at most one item can be produced, or a setup is made ("all or nothing production"). An optimal production schedule for the DLSP minimizes the sum of setup and holding costs. The relationship between the DLSP and scheduling models in general motivated us to solve the DLSP as a batch sequencing problem (BSP). We derive BSP instances from DLSP instances and solve the DLSP as a BSP. Demand for an item is interpreted as a job with a deadline and a processing time. Jobs corresponding to demand for the same item are grouped into one family. Items in the DLSP are families in the BSP. All jobs must be processed on a single machine between time zero and their respective deadlines, while switching from a job in one family to a job in another family incurs (sequence dependent) setup times and setup costs. Early completion of jobs is penalized by earliness costs which correspond to holding costs. As for the DLSP, an optimal schedule for the BSP minimizes the sum of setup costs and earliness costs. The DLSP was first introduced by Lasdon and Terjung [10] with an application to production scheduling in a tire company. Complexity results for the DLSP and its extensions are examined in Salomon et al. [14], where the close relationship of the DLSP to job (class) scheduling problems is emphasized. A broader view on lotsizing and scheduling problems is given in Potts and Van Wassenhove [13]. An approach based on lagrangean relaxation is proposed by Fleischmann [6] for the DLSP without setup times. Fleischmann [7] utilizes ideas from solution procedures for vehicle routing problems to solve the DLSP with sequence dependent setup costs. The DLSP with sequence independent setup times and setup costs is examined by Cattrysse et al. [4]. In a recent work, Salomon et al. [15] propose a dynamic programming based approach for solving the DLSP with sequence dependent setup times and setup costs to optimality. The results of [4], [7] and [15] will serve as a benchmark for our approach for solving the BSP. The complexity of scheduling problems with batch setup times is investigated by Bruno and Downey [2] and Monma and Potts [12]. Bruno and Downey show the feasibility problem to be NP-hard if setup times are nonzero. Solution procedures for scheduling problems with batch setup times are studied in Unal and Kiran [17], Ahn and Hyun [1] and Mason and Anderson [11]. In [17] the feasibility problem of the BSP is addressed and an effective heuristic is proposed. In [1] and [11], algorithms to minimize mean flow time are proposed. Webster and Baker [20] survey recent results and derive properties of optimal schedules for various batching problems. The contribution of the paper is twofold. First, we solve the DLSP as a BSP and state the equivalence between both models such that we can solve either the DLSP or the BSP. Second, we present an algorithm that solves the BSP faster than known procedures solving the DLSP. The paper is organized as follows: we present the DLSP and the BSP in Section 2 and provide a numerical example in Section 3. The relationship of both models is analyzed in Section 4. Section 5 presents a timetabling procedure to convert a sequence into a minimum cost schedule, and in Section 6 we describe a branch & bound algorithm for solving the BSP. A comparison of our algorithm with solution procedures solving variants of the DLSP is found in Section 7. Summary and conclusions follow in Section 8. Model Formulations The DLSP is presented with sequence dependent setup times and setup costs, we refer to this problem as SDSTSC. SDSTSC includes the DLSP with sequence independent setups (SISTSC), sequence dependent setup costs but zero setup times (SDSC), and the generic DLSP with sequence independent setup costs and zero setup times (cf. Fleischmann [6]) as special cases. Table 1: Parameters of the DLSP i index of item (=family), denotes the idle machine t index of periods, q i;t demand of item i in period t holding costs per unit and period of item i st g;i setup time from item g to item i, g;i setup costs per setup period from item g to item i, sc g;i setup costs from item g to item i, st g;i g Table 2: Decision Variables of the DLSP Y i;t 1, if item i is produced in period t, and 0 otherwise. Y time in period t if the machine is setup for item i in period t, while the previous item was item g, and 0 otherwise I i;t inventory of item i at the end of period t The DLSP parameters are given in Table 1. Items (families) in the DLSP (BSP) are indexed by i, and holding costs per unit of item i and period. Production has to fulfill the demand q i;t for item i in period t. Setup costs sc g;i are "distributed" over maxf1; st g;i g setup periods by defining per-period setup costs sc p g;i . The decision variables are given in Table 2: we set Y production takes place for item i in period t. V a setup from item g to item i in period t, and I i;t denotes the inventory of item i at the end of period t. In the mixed binary formulation of Table 3, the objective (1) minimizes the sum of setup costs sc p (per setup period st g;i ) and inventory holding costs. Constraints (2) express the inventory balance. The "all or nothing production" is enforced by constraints (3): in each period, the machine either produces at full unit capacity, undergoes setup for an item, or is idle, i.e. Y for an idle period. For st instantiate V g;i;t appropriately. Constraints (5) couple setup and production whenever st g;i ? 0: if item i is produced in period t and item g in period t\Gamma- \Gamma1 then the decision variable V st g;i . Constraints (6) enforce the correct length of the string of setup variables V g;i;t\Gamma- for st g;i ? 1. However, for st g;i ? 0 we have to exclude the case Y setting any V this is done by constraints (7). Constraints (8) prevent any back-logging. Finally, the variables I i;- , V g;i;- , and Y i;- are initialized for - 0, by constraints (11). Due to the "all or nothing production", we can write down a DLSP schedule in terms of a period-item assignment in a string - specifies the action in each period, i.e. - st g;i ? 0). The BSP is a family scheduling problem (cf. e.g. Webster and Baker [20]). Parameters (cf. Table related to the N families are the index i, the number of jobs n i in each family, and the total number of jobs J . Table 3: Model of the DLSP Min (1) subject to I st st st g;i st g;i ? st g;i st g;i ? 0; I I Table 4: Parameters of the BSP number of jobs of family i, number of jobs denotes the j-th job of family processing time for the j-th job of family i d (i;j) deadline for the j-th job of family i w (i;j) earliness weight per unit time for the j-th job of family i number Table 5: Decision variables of the BSP - sequence of all jobs, denotes the job at position k C (i;j) completion time of job (i; there is idle time between the jobs (i and 0, otherwise Table Model of the BSP Min ZBSP subject to st i [k\Gamma1] ;i [k] (st 0;i [k] st i [k\Gamma1] ;i [k] st i [k\Gamma1] ;i [k] one unit of family i in inventory for one period of time. Setup times st g;i and setup costs sc g;i are given for each pair of families g and i. The set of jobs is partitioned into families i, the j-th job of family i is indexed by the tuple (i; j). Associated with each job (i; are a processing time p (i;j) , a deadline d (i;j) , and a weight w (i;j) . Job weights w (i;j) are proportional to the quantity (=processing time) of the job (proportional weights), they are derived from h i and p (i;j) . We put the tuple in brackets to index the job attributes because the tuple denotes a job as one entity. The decision variables are given in Table 5. The sequence - denotes the processing order of the jobs, denotes the job at position k; together with completion times C (i;j) of each job we obtain the schedule oe. A conceptual model formulation for the BSP is presented in Table 6. ZBSP (oe) denotes the sum of earliness and setup costs for a schedule oe, which is minimized by the objective (12). The earliness weighted by w (i;j) , and setup costs sc i [k\Gamma1] ;i [k] are incurred between jobs of different families. Each job is to be scheduled between time zero and its deadline, while respecting the sequence on the machine as well as the setup times. This is done by constraints (13). Constraints to one if there is idle time between two consecutive jobs. We then have a setup time st 0;i [k] from the idle machine rather than a sequence dependent setup time st i [k\Gamma1] ;i [k] . Initializations of beginning and end of the schedule are given in (16) and (17), respectively. Remark 1 For the BSP and DLSP parameters we assume that: 1. setup times and setup costs satisfy the triangle inequality, i.e. st g;i - st g;l st l;i and sc g;i - 2. there are no setups within a family, i.e. st no tear-down times and costs, i.e. st 3. there is binary demand in the DLSP, i.e. q i;t 2 f0; 1g. 4. jobs of one family are labeled in order of increasing deadlines, and deadlines do not interfere, i.e. Remark 1 states in (1.) that it is not beneficial to perform two setups in order to accomplish one. Mason and Anderson [11] show that problems with nonzero tear-downs can easily be converted into problems with sequence dependent setups and zero tear-downs, which motivates (2. With (1.) and (2.) we have st 0;i - st g;i analogously for setup costs. Thus, the third term in the objective (12) is always nonnegative. Assumption (3.) anticipates the "all or nothing production" for each item i (cf. also Salomon et al. [14]) and is basically the same assumption as (4.): if only jobs of one family i are considered, they can be scheduled with C The main observation that motivated us to consider the DLSP as a special case of the BSP is that the (q i;t )-matrix is sparse, especially if setup times are significant. The basic idea is to interpret items in the DLSP as families in the BSP and to regard nonzero demand in the DLSP as jobs with a deadline and a processing time in the BSP. In order to solve the DLSP as a special case of the BSP we derive BSP instances from DLSP instances in the following way: setup times and setup costs in the BSP and Comparison Equivalence BSP solution procedure DLSP instances DLSP solution procedures BSP(DLSP) or BSPUT(DLSP) Solution Transformation Figure 1: Comparison of DLSP and BSP DLSP are identical, and the job attributes of the BSP instances are derived from the (q i;t )-matrix by Definitions 1 and 2. defined as a BSP instance with unit time jobs derived from a DLSP instance. For each family i there are n jobs. An entry q denotes a job (i; defined as a BSP instance derived from a DLSP instance. A sequence of consecutive "ones" in the (q i;t )-matrix, i.e. q . The number of times that a sequence of consecutive ones appears for an item i defines n i . Figure 1 provides the framework for the BSP-DLSP comparison: after transforming DLSP instances into BSP instances, we compare the performance of solution procedures and the quality of the solutions. The difference between the approaches is as follows: in the DLSP, decisions are made anew in each individual period t, represented by decision variables Y i;t and V g;i;t (cf. Table 2). In the BSP, we decide how to schedule jobs, i.e. we decide about the completion times of the jobs. BSP and DLSP address the same underlying planning problem, but use different decision variables. Br-uggemann and Jahnke [3] make another observation which concerns the transformation of instances: a DLSP instance may be not polynomially bounded in size while the size of the BSP(DLSP) instance is polynomially bounded. On that account, in [3] it is argued, that the (q i;t )-matrix is not a "reasonable" encoding for a DLSP instance in the sense of Garey and Johnson [9] because BSP(DLSP) describes a problem instance in a more concise way. 3 Numerical Example In this section, we provide an example illustrating the generation of BSPUT(DLSP) and BSP(DLSP). We will also refer to this example to demonstrate certain properties of the BSP. In Figure 2 we illustrate the equivalence between both models. The corresponding parameters setup times, setup costs and holding costs are given in Table 7. Figure 2 shows the demand matrix (q i;t ) of DLSP and the jobs at their respective deadlines of BSPUT(DLSP) and BSP(DLSP). Table 7: Numerical Example: Setup and Holding Costs st a 3 a 3 3 3 a 2 1 a a a 1 a 2 BSPUT(DLSP), oe ut a BSPUT(DLSP), oe ut BSP(DLSP), oe d Figure 2: DLSP, BSPUT(DLSP) and BSP(DLSP) Table 8: BSPUT(DLSP) Instance and Solution a For BSPUT(DLSP), we interpret each entry of "one" as a job (i; j) with a deadline d (i;j) . Processing times p (i;j) are equal to one for all jobs. We summarize the BSPUT(DLSP) parameters in Table 8. An optimal DLSP schedule with h is the string - a in Figure 2 (with entries f0; a; 1; 2; 3g for idle or time or for production of the different items, respectively). This schedule is represented by oe ut a for BSPUT(DLSP), and is displayed in Table 8. Both schedules have an optimal objective function value of ZBSP (oe ut a In BSP(DLSP), consecutive "ones" in the demand matrix (q i;t ) are linked to one job. The number of jobs is thus smaller in BSP(DLSP) than in BSPUT(DLSP). For instance, jobs (1; 2) and (1; 3) in BSPUT(DLSP) are linked to one job (1; 2) in BSP(DLSP), compare oe ut b and oe b . However, a BSP(DLSP) schedule cannot a in Figure 2 since there we need unit time jobs. For BSP(DLSP) we now let the cost parameters costs for all families) and sc is the optimal DLSP schedule and oe b the optimal BSP(DLSP) schedule. Again, the optimal objective function value is ZBSP (oe b The example shows that the same schedules can be obtained from different models. In the next section we formally analyze the equivalence between the DLSP and the BSP. 4 Relationship Between BSPUT(DLSP), BSP(DLSP) and DLSP In the BSP we distinguish between sequence and schedule. A BSP schedule may have inserted idle time so that the processing order does not (fully) describe a schedule. In the following we will say that consecutively sequenced before job (i is sequenced at the next position. If we consecutively schedule job (i there is no idle time between both jobs, i.e. the term in brackets in constraints (14) equals zero. A sequence - for the BSP consists of groups, where a group is an (ordered) set of consecutively sequenced jobs which belong to the same family. On the other hand, a schedule consists of (one or several) blocks. Jobs in one block are consecutively scheduled, different blocks are separated by idle time (to distinguish from setup time). Jobs in one block may belong to different families, and both block and group may consist of a single job. As an example refer to Figure 2 where both oe c and oe d consist of five groups, oe c forms two blocks, and oe d is only one block. For a given sequence -, a BSP schedule - oe is called semiactive if C (i;j) is constrained by either d (i;j) or the start of the next job; no job can be scheduled later or rightshifted in a semiactive schedule - oe. We can derive - oe from a sequence - if constraints (13) are equalities and P k is set to zero. The costs ZBSP (-oe) are a lower bound for costs ZBSP (oe) of a BSP schedule oe because - oe is the optimal schedule of the relaxed BSP in which constraints (14) are omitted. However, in the semiactive schedule there may be idle time and it may be beneficial to schedule some jobs earlier, i.e. to leftshift some jobs to save setups (which will be our concern in the timetabling procedure in Section 5). In both models we save setups by batching jobs. In the DLSP, a batch is a non-interrupted sequence of periods where production takes place for the same item i 6= 0, i.e. Y . In the BSP, jobs of one group which are consecutively scheduled without a setup are in the same batch. A batch must not be preempted by idle time. In Figure 2, the group of family 3 forms two batches in schedule oe c whereas this group is one batch in oe b . We will call a sequence - (schedule oe) an EDDWF sequence (schedule) if jobs of one family are sequenced (scheduled) in nondecreasing order of their deadlines (where EDDWF abbreviates earliest deadline within families). Ordering the jobs in EDDWF is called ordered batch scheduling problem in Monma and Potts [12]. By considering only EDDWF sequences, we reduce the search space for the branch & bound algorithm described in Section 6. We first consider BSPUT(DLSP) instances. The following theorem states, that for BSPUT(DLSP) we can restrict ourselves to EDDWF sequences. Theorem 1 Any BSPUT(DLSP) schedule oe can be converted into an EDDWF schedule ~ oe with the same cost. Proof: Recall that jobs of one family all have the same weights and processing times. In a schedule oe, let A; B; C represent parts of oe (consisting of several jobs), and (processing) times of the parts. Consider a schedule where jobs are not ordered in EDDWF, i.e. . The schedule ~ has the same objective function value because w (i;j 1 . The completion times of the parts A; B; C do not change because Interchanging jobs can be repeated until oe is an EDDWF schedule, completing the proof. 2 A DLSP schedule - and a BSPUT(DLSP) schedule oe are called corresponding solutions if they define the same decision. A schedule schedule oe are corresponding solutions if for each point in time the following holds: (i) - in oe the job being processed at t belongs to family a and a setup is performed in oe, and (iii) - and the machine is idle in oe. Figure 2 gives an example for corresponding solutions: - a corresponds to oe ut a , and - b corresponds to oe ut b . We can always derive entries in - from oe, and completion times in oe can always be derived from - if oe is an EDDWF schedule. Theorem 2 A schedule oe is feasible for BSPUT(DLSP) if and only if the corresponding solution - is feasible for DLSP, and - and oe have the same objective function value. Proof: We first prove that the constraints of DLSP and BSPUT(DLSP) define the same solution space. In the DLSP, constraints (2) and (8) stipulate that . For each q i;t ? 0 (2) and (8) enforce a Y t. The sequence on the machine - the sequence dependent setup times taken into account - is described by constraints (3) to (7). In the BSP this is achieved by constraints (13). We schedule each job between time zero and its deadline. All jobs are processed on a single machine, taking into account sequence dependent setup times. Second, we prove that the objective functions (1) and (12) assign the same objective function value to corresponding solutions - and oe: the cumulated inventory for an item i (over the planning horizon equals the cumulated earliness of family i, and job weights equal the holding costs, i.e. h for BSPUT(DLSP). Thus, the terms I i;t and are equal for corresponding solutions - and oe. Some more explanation is necessary to show that corresponding solutions - and oe have the same setup costs. Consider a setup from family g to i (g; i 6= 0) without idle time in oe: we then have sc st g;i g and we have st g;i consecutive "ones" in V g;i;t , which is enforced by (6). On the other hand, in the case of inserted idle time we have a setup from the idle machine (enforced by the decision variable P k of the BSP) and there are st 0;i consecutive "ones" in V 0;i;t . Thus, the terms and are equal for corresponding solutions - and oe. Therefore, corresponding solutions - and oe incur the same holding and the same setup costs, which proves the theorem. 2 As a consequence of Theorem 2, a schedule oe is optimal for BSPUT(DLSP) if and only if the corresponding solution - is optimal for DLSP, which constitutes the equivalence between DLSP and BSP for BSP- UT(DLSP) instances. We can thus solve DLSP by solving BSPUT(DLSP). In general, however, the more attractive option will be to solve BSP(DLSP) because the number of jobs is smaller. Definition 3 In a schedule oe, let a production start of family i be the start time of the first job in a batch. Let inventory for family i build between C (i;j) and d (i;j) . The schedule oe is called regenerative if there is no production start for a family i as long as there is still inventory for family i. The term "regenerative" stems from the regeneration property found by Wagner and Whitin [19] (for similar ideas cf. e.g. Vickson et al. [18]). Each regenerative schedule is also an EDDWF schedule, but the reverse is not true. If a schedule oe is regenerative, jobs (i; are in the same batch if holds. Furthermore, in a regenerative BSPUT(DLSP) schedule oe, jobs from consecutive "ones" in (q i;t ) are scheduled consecutively (recall for instance oe ut b and oe b in Figure 2); hence a regenerative BSPUT(DLSP) schedule represents a BSP(DLSP) schedule as well. In Figure 2, schedule oe d is not regenerative: a batch for family is started at there is still inventory for We first show that we do not lose feasibility when restricting ourselves to regenerative schedules only. Theorem 3 If oe is a feasible BSPUT(DLSP) or BSP(DLSP) schedule then there is also a feasible regenerative schedule ~ oe. A oe ~ oe A Figure 3: Regenerative Schedule Proof: In a schedule oe, let i B (i A ) be the family to which the first (last) job in part B (A) belongs. Consider a non-regenerative schedule oe, i.e. are not in one batch though C Consider schedule ~ where (i; are interchanged and (i; are in one batch. ~ oe is feasible because ~ leftshifting B we do not violate feasibility. Furthermore, due to the triangle inequality we have st i A ;i B - st i A ;i st i;i B . Thus, B can be leftshifted by p (i;j) time units without affecting CA . Interchanging jobs can be repeated until oe is regenerative which proves the theorem. 2 An illustration for the construction of regenerative schedules is depicted in Figure 3. Interchanging (i; and B, we obtain from oe the regenerative schedule ~ oe. Unit processing times are not needed for the proof of Theorem 3, so we have in fact two results: first, to find a feasible schedule we may consider BSP(DLSP) instead of BSPUT(DLSP). Second, for BSP(DLSP) we only need to search over regenerative schedules to find a feasible schedule. Theorem 3 is a stronger result than the one found by Salomon et al. [14] and Unal and Kiran [17] who only state the first result. Moreover, if holding costs are equal, the next theorem extends this result to optimal schedules. Theorem 4 If oe is an optimal BSPUT(DLSP) or BSP(DLSP) schedule and h i is constant 8i, then there is also an optimal regenerative schedule ~ oe. Proof: Analogous to the proof of Theorem 3 we now must consider the change of the objective function value if (i; are interchanged. Without loss of generality, let h ZBSP (oe) (ZBSP (~oe)) denote the objective function value of oe (~oe). For part B, which is leftshifted, we have wB - pB because processing time in part B is at most pB , but B may contain setups as well. Interchanging B and (i; j), the objective changes as follows: Due to the triangle inequality, setup costs and setup times in oe are not larger than in ~ oe, i.e. \Gammasc i A ;i \Gamma explains (i). We leftshift B by p (i;j) and rightshift (i; j) by pB with wB - pB , which explains (ii). Thus ZBSP (~oe) - ZBSP (oe), which proves the theorem. 2 Considering regenerative schedules, we again achieve a considerable reduction of the search space. To summarize we have so far obtained the following results: 1. DLSP and BSP are equivalent for BSPUT(DLSP). 2. Feasibility of BSP(DLSP) implies feasibility of DLSP. 3. For equal holding costs an optimal BSP(DLSP) schedule is optimal for DLSP. When instances with unequal holding costs are solved, the theoretical difference between BSP(DLSP) and DLSP in 3. has only a small effect: computational results in Section 7.3 will show that there is almost always an optimal regenerative BSPUT(DLSP) schedule to be found by solving BSP(DLSP). 5 A Timetabling Procedure for a Given Sequence For a given sequence - the following timetabling procedure decides how to partition - into blocks, or equivalently, which consecutively sequenced jobs should be consecutively scheduled. In the BSP model formulation of Table 6, we have the job at position k starts a new block, or P blocked with the preceding job. By starting a new block at position k, we save earliness costs at the expense of additional setup costs. In Figure 2, earliness costs of oe b are higher than for oe c but we save one setup in oe b . In the timetabling procedure, we start with the semiactive schedule and leftshift some of the jobs to find a minimum cost schedule. Consider the example in Figure 2: for the sequence (3; 3); (2; 2); (1; 2)) the semiactive schedule - oe is given in Figure 4. We first consider two special cases. If we omit constraints (14) of the BSP (so that each group is a batch and idle time may preempt the batch) timetabling is trivial: the semiactive schedule is optimal for a given sequence because no job can be rightshifted to decrease earliness costs (and because setup costs are determined by - and not by oe). Timetabling is also trivial if earliness weights are zero (i.e. h case, we can leftshift each job (without increasing earliness costs) until the resulting schedule is one block (e.g. schedule oe d in Figure 2 is one block and no job can be leftshifted). We have sc g;i - sc i;0 +sc setup costs are minimized if jobs are scheduled in a block, and there is an optimal schedule which consists of one block. In the general case, we need some definitions: block costs bc k1;k2 are the cost contribution of a block from position k 1 to k 2 , i.e. bc k1 The block size bs k is the number of jobs which are consecutively scheduled after job (i included). For instance, in Figure 4 we have bs Let denote f k (b) the costs of a schedule from position k to J if bs k the costs of the minimum cost schedule and bs k the corresponding block size at position k. The recurrence equation for determining f k and bs f b=1;:::;bs oe Position k Figure 4: Semiactive Schedule Table 9: Computations of Equation (18) for the Example in Figure 4 f In equation (18) we take the minimum cost for bs is the maximum block size at position k (and a new block starts at k+bs 1). For a given block size b, f k (b) is the sum of block costs from position k to position k to the next block and the minimum cost f k+b . Basically, equation (18) must be computed for every sequence. However, some simplifications are possible: if two jobs can be consecutively scheduled in the semiactive schedule, it is optimal to increment bs bs so that equation (18) needs not to be evaluated. Consequently, if the semiactive schedule is one block, timetabling is again trivial: each group in - oe equals one batch, the whole schedule forms a block, and - If setups are sequence independent, a minimum cost schedule can be derived with less effort as follows: let the group size gs k at position k denote the number of consecutively sequenced jobs at positions r ? k that belong to the same family as job (i [k] ; j [k] ). Then, for sequence independent setups, equation (18) must be evaluated only for gs k . The reasoning is as follows: jobs of different groups are leftshifted to be blocked only if we can save setup cost. Then, for consecutive groups of families g and i we would need sc which does not hold for sequence independent setups; therefore, we only need to decide about the leftshift within a group. An example for the computations of equation (18) is given in Table 9 for the semiactive schedule in Figure 4 (see the cost parameters in Table 7). The schedule - oe contains idle time, and we determine f and bs k for each position k, k - J . Consider jobs (2,2), (3,3) and (3,2) at positions 6,5 and 4 in Figure 4. Up to position 4 the semiactive schedule is one block, and we increment bs k , which is denoted by entries (-) in Table 9. After job (3,1), oe has inserted idle time between positions 3 and 4 (d different block sizes must be considered to find the minimum cost schedule. In Table 9, we find f i.e. we start a new block after position split the group into two batches, as done for oe c in Figure 2. For the objective function value, we add sc 0;2 to f 1 and obtain a cost of ZBSP (oe c 6 Sequencing Algorithm In this section we present a branch & bound algorithm for solving the BSP to optimality, denoted as SABSP. Jobs are sequenced backwards, i.e. at stage 1 a job is assigned to position J , at stage 2 to position stage s to position assigns s jobs to the last s positions of sequence -, in addition to that an s-partial schedule oe s also assigns completion times to each job in - s . A partial schedule ! s is called completion of oe s if ! s extends oe s to a schedule oe which schedules all jobs, and we write We only examine EDDWF sequences, as if there were precedence constraints between the jobs. The precedence graph for the example in Figure 2 is shown in Figure 5. Using the EDDWF ordering, we Figure 5: EDDWF Precedence Graph for Backward Sequencing Table 10: Attributes of Partial Schedules under consideration at stage s UB upper bound, objective function value of the current best schedule c(oe s ) cost of oe s without the setup for (i s AS s (US s ) set of jobs already scheduled (unscheduled) in the s-partial schedule oe s UI s set of families to which jobs in US s belong to of jobs which form the first block of oe s of earliness weights of jobs in G 1 (oe s ), i.e. decide in fact at each stage s which family to schedule. A job is eligible at stage s if all its (precedence related) predecessors are scheduled. An s-partial schedule (corresponding to a node in the search tree) is extended by scheduling an eligible job at stage s + 1. We apply depth-first search in our enumeration and use the bounding, branching, and dominance rules described in Sections 6.1 and 6.2 to prune the search tree. Each (s-partial) sequence - s uniquely defines a minimum cost (s-partial) schedule oe s by the timetabling procedure. Enumeration is done over all sequences and stops after all sequences have been (implicitly) examined; the best solution found is optimal. The implementation of SABSP takes advantage of the fact that equation (18) needs not be recalculated for every oe s and, in the case of backtracking, the computation of equation (18) has already been accomplished for the partial schedule to which we backtrack. Table lists attributes of s-partial schedules. For each scheduling stage s we identify the job (i s under consideration, and the start time t(oe s ) and costs c(oe s ) of the s-partial schedule. The set of currently scheduled (unscheduled) jobs is denoted by AS s (US s ). UI s denotes to which families the jobs in US s belong; UB is the current upper bound. 6.1 Bounding and Branching Rules The feasibility bound states that for a given oe s , all currently unscheduled jobs in US s must be scheduled between time zero and t(oe s ) and, furthermore, we need a setup time for each family in UI s . More formally, define min fst (i;j)2US s Then, oe s has no feasible completion ! s if t(oe s The cost bound states that costs c(oe s ) of an s-partial schedule are a lower bound for all extensions of oe s , and for any completion ! s at least one setup for each family in UI s must be performed. We define min fsc g;i g: Then, oe s cannot be extended to a schedule improves UB if C s Both bounds are checked for each s-partial schedule oe s . Clearly, T s and C s can be easily updated during the search. We also tested a more sophisticated lower bound where all unscheduled jobs were scheduled in EDD order without setups. In this way we were able to derive a lower bound on the earliness costs as well and check feasibility more carefully, but computation times did not decrease. If only regenerative schedules need to be considered to find the optimal schedule (cf. Theorem 4), we employ a branching rule as follows: scheduling (i s eligible in the EDDWF precedence graph. If t(oe s we batch (i s and do not consider any other job as an extension of oe s . We need not enumerate partial schedules where oe s is extended by a job (g; because then the resulting schedule is non-regenerative. 6.2 Dominance Rules The most remarkable reduction of computation times comes as a result of the dominance rules. The dominance rules of SABSP compare two s-partial schedules oe s and oe s , which schedule the same set of jobs, so that AS s . In this notation, schedule oe s denotes the s-partial schedule currently under consideration, while oe s denotes a previously enumerated schedule which may dominate oe s . A partial schedule oe s dominates oe s if it is more efficient in terms of time and cost: oe s starts later to schedule the job-set, i.e. t(oe s incurs less cost, i.e. c(oe s ) - c(oe s ). If the family i s of the job scheduled at stage s differs in oe s and oe s we make both partial schedules "comparable" with a setup from i s to we compare time and cost but subtract setup times and setup costs appropriately. If a schedule oe s is not dominated, we store for the job set AS s and family i s the pair t(oe s ) and c(oe s ) which is "most likely" to dominate other s-partial schedules. Note that the number of partial schedules is exponential in the number of items N so that storage requirements for the dominance rules grow rapidly if N increases. For a formal description of the dominance rules we need several definitions (cf. Table 10): all jobs which form a block with (i s belong to the set G 1 (oe s ), and the sum of earliness weights in G 1 (oe s ) is denoted as w 1 (oe s ). The dominance rules take into account the block costs for all extensions of oe s and oe s : we consider for oe s the maximum, for oe s the minimum costs incurred by blocking; oe s then dominates oe s if c(oe s ) plus an upper bound on block costs is less or equal c(oe s ) plus a lower bound on block costs. An upper bound on the block costs for oe s is given by sc 0;i s (recall that sc 0;i - sc g;i ). Then, oe s starts a new block. But a tighter upper bound can be found for start times close to t(oe s in order to save costs we can leftshift all the jobs in G 1 (oe s ) (but only these), because after G 1 (oe s ) we perform a new setup from the idle machine. G 1 (oe s ) is the largest block which may be leftshifted. Let pbt(oe s ) denote the time where the cost increase due to a leftshift of G 1 (oe s ) exceeds sc 0;i s . We then have w 1 (oe s )(t(oe s and define the pull-back-time pbt(oe s ) of an s-partial schedule oe s as follows: Consequently, for time t, pbt(oe s an upper bound on block costs is given by leftshifting G 1 (oe s ); are bounded by sc 0;i s . A lower bound on the block costs for oe s is given in the same way as for oe s , but now we consider the smallest block that can be leftshifted, which is simply job (i s We can now state the dominance rule: we differentiate between i (Theorem 5) and i s orem 6). Theorem 5 Consider two s-partial schedules oe s and oe s with AS Proof: Any completion ! s of oe s is also a feasible completion of oe s because of (i); if (! s ; oe s ) is feasible, too. Due to (ii), for any ! s , the schedule (! s ; oe s ) has lower costs than (! s ; oe s ). In the following we consider the cost contributions of oe s and oe s due to leftshifting, when we extend oe s and oe s . Consider Figure 6 for an illustration of the situation in a time-cost diagram. We have due to EDDWF also (i s line represents the upper bound on block costs for oe s . For pbt(oe s expensive to leftshift G 1 (oe s ), while for t ! pbt(oe s ) a setup from the idle machine to i s is performed. The broken line represents the lower bound on block costs for oe s . The smallest block that can be leftshifted is the job (i s In order to prove that oe s will never have less costs than oe s due to blocking, we check the costs at points (ii) and (iii): at (ii) we compare the costs at t(oe s ) while at (iii) we compare them at pbt(oe s ). Between (ii) and (iii) costs increase linearly, and for t ! pbt(oe s ) we know that there is a monotonous cost increase for oe s , while costs of oe s no longer increase. Thus, if (ii) and (iii) are fulfilled, cost contributions of oe s are less than those of oe s , i.e. there is no completion ! s such that ZBSP (! s ; oe s completing the proof. 2 For the example in Figure 2, Figures 7 and 8 illustrate Theorem 5 with 3-partial schedules oe 3 and oe 3 . In Figure 7, we have G 1 (oe 3 cost oe s oe s Figure Illustration of Theorem 5 oe 3t Figure 7: Theorem 5: oe 3 dominates oe 3 Checking (ii), we have 22 - while for (iii) we have 22 so that oe 3 is dominated. Figure 8 illustrates the effect of block costs, but with modified data as follows: 10=3. Checking Theorem 5, we have (ii) not fulfilled. Thus, oe 3 does not dominate oe 3 , though c(oe 3 Figure 8 shows that c(oe 4 leftshifted. In the second dominance rule for the case i s must consider sc 0;i instead of sc g;i to take block costs into account. Theorem 6 Given two s-partial schedules oe s and oe s with AS st denote the family of the (last) job in a completion ! s of oe s . Then st st st analogously for setup costs due to the triangle inequality. Thus any completion ! s of oe s is also a feasible completion of oe s because of (i); if (! s ; oe s ) is feasible, (! s ; oe s ) is feasible, too. Due to (ii), for any ! s , the schedule (! s ; oe s ) has lower costs than (! s ; oe s ). Figure 8: Theorem 5: oe 3 does not dominate oe 3 The difference is that now also block costs are taken into account in (ii): when leftshifting G 1 (oe s ) in an extension of oe s , we have c(oe s as an upper bound for the cost contribution. A trivial lower bound for the cost contribution of oe s is c(oe s ). Thus oe s dominates oe s as any ! s completes oe s at lower costs, completing the proof. 2 Finally, an alternative way to solve the BSP is a dynamic programming approach. We define the job-sets as states and apply the dominance rules in the same way. An implementation of this approach was less efficient and is described in Jordan [8]. 7 Comparison with Procedures to solve Variants of the DLSP From the analysis in Section 4 we know that we address the same planning problem in BSP and DLSP, and that we find corresponding solutions. Consequently, in this section we compare the performance of algorithms solving the BSP with procedures for solving variants of the DLSP. The comparison is made on the DLSP instances used to test the DLSP procedures; we take the instances provided by the cited authors and solve them as BSP(DLSP) or BSPUT(DLSP) instances (cf. Figure 1). An exception is made for reference [7] where we use randomly generated instances. The different DLSP variants are summarized in Table 11. For the DLSP, in the first column the reference, in the second the DLSP variant is displayed. The fourth column denotes the proposed algorithm, the third column shows whether computational results for the proposed algorithm are reported for equal or unequal holding costs. Depending on the holding costs, the different DLSP variants are solved as BSP(DLSP) or BSPUT(DLSP) instances. With the exception of reference [15], the DLSP procedures are tested with equal holding costs, so that regenerative schedules are optimal in [4] and [7]. 7.1 Sequence Independent Setup Times and Setup Costs (SISTSC) In Cattrysse et al. [4], a mathematical programming based procedure to solve SISTSC is proposed. Cattrysse et al. [4] refer to their procedure as dual ascent and column generation procedure (DACGP). The DLSP is first formulated as a set partitioning problem (SPP) where the columns represent the production schedule for one item i; the costs of each column can be calculated separately because setups are sequence independent. DACGP then computes a lower bound for the SPP by column generation, new Table Solving Different DLSP Variants as a BSP Author Variant Holding Costs Algorithm Instances Properties of Schedules Cattrysse et al. [4] regenerative Fleisch- mann regenerative Salomon et al. [15] one block columns can be generated solving a single item subproblem by a (polynomial) DP recursion. In DACGP a feasible schedule, i.e. an upper bound, may be found in the column generation step, or is calculated by an enumerative algorithm with the columns generated so far. If in neither case a feasible schedule is found, an attempt is made with a simplex based procedure. The (heuristic) DACGP generates an upper and a lower bound, SABSP solves BSP(DLSP) to optimality. DACGP is coded in FORTRAN, SABSP is coded in C. DACGP was run on an IBM-PS2 Model 80 PC (80386 processor) with a 80387 mathematical coprocessor, we implemented SABSP on the same machine to make computation times comparable. Computational results for the DACGP are reported only for identical holding costs items. Consequently, we solve DLSP as BSP(DLSP) and only need to consider regenerative schedules, Theorem 4. Furthermore, the timetabling procedure requires fewer computations in equation (18) as setups are sequence independent. The DLSP instances with nonzero setup times are provided by the authors of [4]. They generated instances for item-period combinations f(N; T 60)g. We refer only to instances with smaller instances are solved much faster by SABSP than by DACGP. The DLSP instances have setup times st g;i of either 0, 1 or 2 periods. The average setup-time per item (over all instances) is (approximately) 0.5, making setup times not very significant. For each item- period combination instances with different (approximate) capacity utilizations ae were generated: low (L) capacitated (ae ! 0:55), medium (M) (0:55 - ae ! 0:75) and high (H) capacitated instances (ae - 0:75). Approximate capacity utilization is defined as ae = 1=T were generated for each combination, amounting to 3 instances in total. In Table 12, we use #J to denote the average number of jobs in BSP(DLSP) for the instance size (N; T ) of the DLSP. For DACGP we use 4 avg to denote the average gap (in percent) between upper and lower bound. # inf is the number of instances found infeasible by the different procedures and R avg denotes the average time (in seconds) needed for the instances in each class. For DACGP, all values in Table 12 are taken from [4]. Table 12: Comparison of DLSP and BSP Algorithms for SISTSC DACGP SABSP (4; (386 PC with coprocessor) In the comparison between DACGP and SABSP, the B&B algorithm solves problems with much faster; the number of sequences to examine is relatively small. For computation times of SABSP are in the same order of magnitude than for DACGP. In (6; 60;M) the simplex based procedure in DACGP finds a feasible integer solution for one of the 10 instances claimed infeasible by DACGP. Thus, in (6; 60;M), 9 instances remain unsolved by DACGP, whereas SABSP finds only 7 infeasible instances. DACGP also fails to find existing feasible schedules for (N; T; ae) =(2,60,H), (4,60,M). Recall that SABSP takes advantage of a small solution space, keeping the enumeration tree small and thus detecting infeasibility or a feasible schedule quite quickly. DACGP tries to improve the lower and upper bound, which is difficult without an initial feasible schedule. Therefore the (heuristic solution procedure) DACGP may fail to detect feasible schedules if the solution space is small. For the same problem size (N; T ) in DLSP, the number of jobs J in BSP(DLSP) may be very different. Therefore, solution times differ considerably for SABSP. Table 13 presents the frequency distribution of solution times. In every problem class the majority of instances is solved in less than the average time for DACGP. 7.2 Sequence Dependent Setup Costs (SDSC) An algorithm for solving SDSC is proposed by Fleischmann [7]. Fleischmann transforms the DLSP into a traveling salesman problem with time windows (TSPTW), where a tour corresponds to a production schedule in SDSC. Fleischmann calculates a lower bound by lagrangean relaxation; the condition that each node is to be visited exactly once, is relaxed. An upper bound is calculated by a heuristic, that first constructs a tour for the TSPTW and then tries to improve the schedule using an Or-opt operation. In Or-opt, pieces of the initial tour are exchanged to obtain an improved schedule. Or-opt is repeated until no more improvements are found. We refer to Fleischmann's algorithm as TSPOROPT. TSPOROPT Table 13: Frequency Distribution of Solution Times of SABSP Number of instances solved faster than (4; 28 (386 PC with coprocessor) was coded in Fortran, experiments were performed on a 486DX2/66 PC with the original code provided by Fleischmann. Fleischmann divides the time axis into micro and macro periods. Holding costs arise only between macro periods, and demand occurs only at the end of macro periods. Thus a direct comparison of TSPOROPT and SABSP using Fleischmann's instances is not viable; instead, we use randomly generated BSP instances which are then transformed into DLSP instances. We generated instances for low (L) (ae - 0:75) or high (H)(ae - 0:97) capacity utilization. Note that for zero setup times, ae does not depend on the schedule; the feasibility problem is polynomially solvable. In BSP, we have an average number of jobs with a processing time out of the interval [1; 4]. In DLSP, we have an average for high (H) and low (L) capacitated instances. Holding costs are identical, and we solve BSP(DLSP). From [7] we select the 2 setup cost matrices S4 and S6 which satisfy the triangle inequality: in S4 costs equal 100 for g ! i and 500 for g ? i. For S6 we have only two kinds of setups: items f1; 2; 3g and f4; 5g form two setup-groups, with minor setup costs of 100 within the setup-groups and major setup costs of 500 from one setup-group to the other. In Table 14 results are aggregated over the instances in each class. We use 4 avg to denote the average gap between lower and upper bound in % for TSPOROPT and R avg ( ~ R avg ) to denote the average time for TSPOROPT (SABSP) in seconds. We denote by 4Z best the average deviation in % of the objective function value of the heuristic TSPOROPT from the optimal one found by SABSP. Table 14 shows that 4 avg can be quite large for TSPOROPT. Solution times of SABSP are short for high capacitated instances and long for low ones. For S4, TSPOROPT generates a very good lower bound, we have best and the deviation from the optimal objective is due to the poor heuristic upper bound. On the other hand, for S6 both the lower and the upper bound are not very close to the optimum. It is well to note that SABSP does not solve large instances of SDSC with 8 or 10 items whereas Fleischmann reports computational experience for instances of this size as well. The feasibility bound is much weaker for zero setup times, or, equivalently, the solution space is much larger, making SABSP less effective. For Table 14: Comparison of DLSP and BSP Algorithms for SDSC setup cost TSPOROPT SABSP best R avg ~ R avg the instances in Table 14, however, SABSP yields a better performance. 7.3 Sequence Dependent Setup Times and Setup Costs (SDSTSC) In Salomon et al. [15], Fleischmann's transformation of the DLSP into a TSP with time windows (TSPTW) is extended for nonzero setup times in order to solve SDSTSC. Nodes in the TSP network represent positive demands, and all nodes must be visited within a certain time window. The transformed DLSP is solved by a dynamic programming approach designed for TSPTW problems (cf. Dumas et al. [5]), we refer to the procedure in [15] as TSPTWA. Paths in the TSP network correspond to partial schedules. Similar to the dominance rule for SABSP, in TSPTWA paths may dominate other paths via a cost dominance, or they may be eliminated because they cannot be extended, which corresponds to the feasibility bound. TSPTWA is coded in C and run on a HP9000/730 workstation (76 mips, 22 M flops). SABSP runs on a 486DX2/66 PC. In order to test TSPTWA Salomon et al. [15] use randomly generated instances, in which, similar to [4], setup times st g;i 2 f0; 2g. Unfortunately, the setup times do not satisfy the triangle inequality. A "triangularization" (e.g. with the Floyd/Warshall algorithm) often results in setup times equal to zero. So we adjusted the setup times "upwards" (which is possible in this case because st g;i 2 f0; 1; 2g) and as a result, setup times are rarely zero. We added 4 (8) units to the planning horizon for in order to obtain the same (medium) capacity utilization as in [15]. In this way, instances are supposed to have the same degree of difficulty for TSPTWA and SABSP: the smaller solution space due to correcting st g;i upwards is compensated by a longer planning horizon. In [15] instances are generated for and we take the (largest) instances for the item-period combination f(N; T 60)g. The instances have a medium (M) capacity utilization 0:5 - ae - 0:75 because setup times are nonzero. For each (N; T ) combination, with and without holding costs are generated. Holding costs differ among the items. Consequently, we solve BSPUT(DLSP) if Furthermore, we need not apply the timetabling procedure in the latter case because the optimal schedule is one block. In Table 15, #F (# ~ the number of problems solved by TSPTWA (SABSP) within a time limit of 1200 sec (1200 sec) and a memory limit of 20 MB (10 MB). #J denotes the average number of jobs for the BSP. ~ R avg ( ~ Table 15: Comparison of DLSP and BSP Algorithms for SDSTSC F ~ R avg ~ denotes the average time SABSP requires to solve the instances (considering only regenerative schedules). The average time is calculated over all instances which are solved within the time limit, ~ R avg is put in brackets if not all instances are solved. The last column shows the results if we consider only regenerative schedules during enumeration for provides the maximal deviation in % from the optimal schedule (which may be non-regenerative). Table 15 demonstrates that SABSP succeeds in solving some of the problems which remained unsolved by TSPTWA. Solution times of SABSP are relatively short compared with TSPTWA for 5. Solution times increase for and instances can only be solved if the number of jobs is relatively small. Instances become difficult for nonzero, especially for unequal holding costs. If we only enumerate over regenerative schedules, solution times for SABSP decrease. Moreover, only one instance is not solved to optimality for (N; T Thus, even for unequal holding costs optimal schedules are regenerative in most cases. Furthermore, for (N; T instances would have been solved within the time limit of 1200 sec if only regenerative schedules would have been considered. 8 Summary and Conclusions In this paper, we examined both the discrete lotsizing and scheduling problem (DLSP) and the batch sequencing problem (BSP). We presented model formulations for the DLSP and for the BSP. In the DLSP, decisions regarding what is to be done are made in each individual period, while in the BSP, we decide how to schedule jobs. The DLSP can be solved as a BSP if the DLSP instances are transformed. For each schedule of one model there is a corresponding solution for the other model. We proved the equivalence of both models, meaning that for an optimal schedule of the BSP the corresponding solution of the DLSP is also an optimal schedule, and vice versa. In order to solve the BSP effectively, we tried to restrict the search to only a subset of all possible schedules. We found out that jobs of one family can be preordered according to their deadlines. Furthermore, for equal holding costs, it is optimal to start production for a family only if there is no inventory of this family. When solving the BSP with a branch & bound algorithm to optimality, we face the difficulty that already the feasibility problem is difficult. We must maintain feasibility and minimize costs at the same time. Compared with other scheduling models, the objective function is rather difficult for the BSP. A tight lower bound could thus not be developed. We therefore used dominance rules to prune the search tree. Again, the difficult objective function complicates the dominance rules and forces us to distinguish different cases. In order to evaluate our approach, we tested it against (specialized) procedures for solving variants of the DLSP. Despite the fact that we have no effective lower bound, our approach proved to be more efficient if (i) the number of items is small, and (ii) instances are hard to solve, i.e. capacity utilization is high and setup times are significant. It is then "more appropriate" to schedule jobs than to decide what to do in each individual period. In the DLSP, the time horizon is divided into small periods and all parameters are based on the period length. In the BSP, all parameters can also be real numbers: setup times, in particular, are not restricted to being multiples of a period length. The different models also result in different problem sizes for DLSP and BSP: the problem size for DLSP is essentially the number of items N and periods T while the problem size for the BSP depends on the number of families and jobs. We conjecture that our approach is advantageous for instances with few items and a small solution space (i.e. long setup times and high capacity utilization), where the job sequence is the main characteristic of a solution. In such cases we managed to solve instances with 10 (5) families and jobs on a PC. DLSP solution procedures are thought to be better suited for lower capacitated instances with many items, setup times that are not very significant, and parameters which differ among the periods. It is then appropriate to decide anew for each individual period. In the future we will extend the BSP to multilevel structures and multiple machines. Acknowledgments We are indebted to Dirk Cattrysse and Marc Solomon who made available their instances, and to Bernhard Fleischmann who made available his code. Furthermore, we would like to thank three anonymous referees for their valuable comments on earlier versions of this paper. --R Single facility multi-class job scheduling Complexity of task sequencing with deadlines "Some extensions of the discrete lotsizing and scheduling problem" A dual ascent and column generation heuristic for the discrete lotsizing and scheduling problem with setup-times Technical Note: An optimal algorithm for the traveling salesman problem with time windows. The discrete lot-sizing and scheduling problem The discrete lot-sizing and scheduling problem with sequence-dependent setup-costs Batching and Scheduling - Models and Methods for Several Problem Classes Computers and intractability - a guide to the theory of NP-completeness Minimizing flow time on a single machine with job classes and setup times. On the complexity of scheduling with batch setup-times Integrating scheduling with batching and lot- sizing: a review of algorithms and complexity Some extensions of the discrete lotsizing and scheduling problem. Discrete lotsizing and scheduling with sequence dependent setup times and costs. Batching in single operation manufacturing systems. Batch sequencing. Batching and sequencing of components at a single facility. Dynamic version of the economic lot size model. Scheduling groups of jobs on a single machine. --TR --CTR C. K. Y. Lin , C. L. Wong , Y. C. Yeung, Heuristic Approaches for a Scheduling Problem in the Plastic Molding Department of an Audio Company, Journal of Heuristics, v.8 n.5, p.515-540, September 2002 Satyaki Ghosh Dastidar , Rakesh Nagi, Scheduling injection molding operations with multiple resource constraints and sequence dependent setup times and costs, Computers and Operations Research, v.32 n.11, p.2987-3005, November 2005

batch sequencing;Sequence-Dependent Setup Times and Setup Costs;Bounding/Dominance Rule;Discrete Lotsizing and Scheduling;Branch-and-Bound Algorithm

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Synthesis of Novel Views from a Single Face Image.

Images formed by a human face change with viewpoint. A new technique is described for synthesizing images of faces from new viewpoints, when only a single 2D image is available. A novel 2D image of a face can be computed without explicitly computing the 3D structure of the head. The technique draws on a single generic 3D model of a human head and on prior knowledge of faces based on example images of other faces seen in different poses. The example images are used to learn a pose-invariant shape and texture description of a new face. The 3D model is used to solve the correspondence problem between images showing faces in different poses. The proposed method is interesting for view independent face recognition tasks as well as for image synthesis problems in areas like teleconferencing and virtualized reality.

Introduction Given only a driver's license photograph of a person's face, can one infer how the face might look like from a different viewpoint? The three-dimensional structure of an object determines how the image of the object changes with a change in viewpoint. With viewpoint changes, some previously visible regions of the object become oc- cluded, while other previously invisible regions become visible. Additionally, the arrangement or configuration of object regions that are visible in both views may change. Accordingly, to synthesize a novel view of an object, two problems must be addressed and resolved. First, the visible regions that the new view shares with the previous view must be redrawn at their new positions. Sec- ond, regions not previously visible from the view of the example image must be generated or syn- thesized. It is obvious that this latter problem is unsolvable without prior assumptions. For human which share a common structure, such prior knowledge can be obtained through extensive experience with other faces. The most direct and general solution for the synthesis of novel views of a face from a single example image is the recovery the three-dimensional structure of the face. This three-dimensional model can be rotated artificially and would give the correct image for the all points visible in the example image (i.e. the one from which the model was obtained). However, without additional assump- tions, the minimal number of images necessary to reconstruct a face using localized features is three (Huang and Lee, 1989), and even the assumption that a face is bilaterally symmetric reduces this number only to two (Rothwell et al., 1993; Vetter and Poggio, 1994). While shape from shading algorithms have been applied in previous work to recover the surface structure of a face (Horn, 1987), the inhomogeneous reflectance properties of faces make surface integration over the whole face imprecise and questionable. Additionally, the fact that the face regions visible from a single image are insufficient to obtain the three-dimensional structure makes clear, that the task of synthesizing new views to a given single image of a face, cannot be solved without prior assumptions about the structure and appearance of faces in general. Models that have been proposed previously to generalize faces from images can be subdivided into two groups: those drawing on the three-dimensional head structure and those considering only view- or image-dependent face models. In general, the knowledge about faces, which has been incorporated into flexible three-dimensional head models, consists of hand-constructed representations of the physical properties of the muscles and the skin of a face (Terzopoulos and Waters, 1993; Thalmann and Thalmann, 1995). To adjust such a model to a particular face, two or more images were used (Akimoto et al., 1993; Aizawa et al., 1989). For present purposes, it is difficult to assess the usefulness of this approach, since generalization performance to new views from a single image only has never been reported. In recent years, two-dimensional image-based face models have been applied for the synthesis of rigid and nonrigid face transitions (Craw and Cameron, 1991; Poggio and Brunelli, 1992; Beymer et al., 1993; Cootes et al., 1995). These models exploit prior knowledge from example images of prototypical faces and work by building flexible image-based representations (active shape models) of known objects by a linear combination of labeled examples. These representations are applied for the task of image search and recognition (Cootes et al., 1995) or synthesis (Craw and Cameron, 1991). The underlying coding of an image of a new object or face is based on linear combinations of the two-dimensional shape of examples of prototypical images. A similar method has been used to synthesize new images of a face with a different expression or a changed viewpoint (Beymer et al., 1993) making use of only a single given image. The power of this technique is that it uses an automated labeling algorithm that computes the correspondence between every pixel in the two images, rather than for only a hand-selected subset of feature points. The same technique has been applied recently to the problem of face recognition across viewpoint change with the aim of generating additional new views given an example face image (Beymer and Poggio, 1995). In spite of the power of this technique, its most serious limitation is its reliance on the solution of the correspondence problem across view changes. Over large changes in viewpoint, this is still highly problematic due to the frequency with which occlusions and occluding contours occur. To overcome this difficulty in the present work, we draw on the concept of linear object classes, which we have introduced recently in the context of object representations (Vetter and Poggio, 1996). The application of the linear object class approach to this problem mediates the requirement of image correspondence across large view changes for success in novel view synthesis. MAPPING PROCESS Figure 1: Two examples of face images (top row) mapped onto a reference face (center) using pixelwise correspondence established through an optical flow algorithm are shown (lower row). This separates the 2D-shape information captured in the correspondence field from the texture information captured in the texture mapped onto the reference face (lower row). Overview of the Approach In the present paper, the linear object class approach is improved and combined with a single three-dimensional model of a human head for generating new views of a face. By using these techniques in tandem, the limitations inherent in each approach (used alone) can be overcome. Specif- ically, the present technique is based on the linear object class method described in (Vetter and Poggio, 1996), but is more powerful because the addition of the 3D model allows a much better utilization of the example images. The 3D-model also allows the transfer of features particular to an individual face from the given example view into new synthetic views. This latter point is an important addition to the linear class approach, because it now allows for individual identifying features like moles and blemishes that are present in "non- standard" locations on a given individual face, to be transferred onto synthesized novel views of the face. This is true even when these blemishes, etc., are unrepresented in the "general experience" that the linear class model has acquired from example faces. On the other hand, the primary limitation of a single 3D head model is the well-known difficulty of representing the variability of head shapes in general, a problem that the linear class model, with its exemplar-based knowledge of faces will allow us to solve. Another way of looking at the combination of these approaches returns us to the two-fold problem we described at the beginning of this paper. The synthesis of novel views from a single exemplar image requires the ability to redraw the regions shared by the two views, and also the ability to generate the regions of the novel face that are invisible in the exemplar view. The 3D head model allows us to solve the former, and linear object class approach the allows us to solve the latter. Linear Object Classes A linear object class is defined as a 3D object class for which the 3D shape can be represented as a linear combination of a sufficiently small number of prototypical objects. Objects that meet this criterion have the following important property. New orthographic views according to uniform affine 3D transformation can be generated for any object of the class. Specifically, rigid transformations in 3D, can be generated exactly if the corresponding transformed views are known for the set of proto- types. Thus, if the training set consists of frontal and rotated views of a set of prototype faces, any rotated view of a new face can be generated from a single frontal view - provided that the linear class assumption holds. The key to this approach is a representation of an object or face view in terms of a shape vector and a texture vector (see also (Cootes et al., 1995; Jones and Poggio, 1995; Beymer and Pog- gio, 1995)). The separation of 2D-shape and texture information in images of human faces requires correspondence to be established for all feature points. At its extreme, correspondence must be established for every pixel, between the given face image and a reference image. As noted previ- ously, while this is an extremely difficult problem when large view changes are involved, the linear object class assumption requires correspondence only within a given viewpoint - specifically, the correspondence between a single view of an individual face and a single reference face imaged from the same view. Separately for each orien- tation, all example face images have to be set in correspondence to the reference face in the same pose, correspondence between different poses is not needed. This can be done off-line manually (Craw and Cameron, 1991; Cootes et al., 1995) or automatically (Beymer et al., 1993; Jones and Poggio, 1995; Beymer and Poggio, 1995; Vetter and Poggio, 1996). Once the correspondence problem within views is solved, the resultant data can be separated into a shape and texture vector. The shape vector codes the 2D-shape of a face image as deformation or correspondence field to a reference face (Beymer et al., 1993; Jones and Poggio, 1995; Beymer and Poggio, 1995; Vetter and Pog- gio, 1996), which later also serves as the origin of a linear vector space. Likewise the texture of the exemplar face is coded in a vector of image intensities being mapped onto corresponding positions in the reference face image (see also figure 1 lower row). The Three-dimensional Head Model The linear class approach works well for features shared by all faces (e.g. eyebrows, nose, mouth or the ears). But, it has limited representational possibilities for features particular to a individual face (e.g. a mole on the cheek). For this reason, a single 3D model of a human head is added to the linear class approach. Face textures mapped onto the 3D model can be transformed into any image showing the model in a new pose. The final "ro- tated" version of a given face image (i.e. including moles, etc.) can be generated by applying to this new image of the 3D model the shape transformation given through the linear object class ap- proach. This is described in more detail shortly. The paper is organized as follows. First, the algorithm for generating new images of a face from a single example image is described. The technical details of the implementation used to realize the algorithm on grey level images of human faces are described in the Appendix. Under Results a comparison of different implementations of the generalization algorithm are shown. Two variations of the combined approach are compared with a method based purely on the linear object class as described previously (Vetter and Poggio, 1996). First, the linear class approach is applied to the parts of a face separately. The individual parts in the two reference face images were separated using the 3D-model. Second, the 3D-model was used additionally to establish pixelwise correspondence between the two reference faces images in the two different orientations. This correspondence field allows texture mapping across the view point change. Finally, the main features and possible future extensions of the technique are discussed Approach and Algorithm In this section an algorithm is developed that allows for the synthesis of novel views of a face from from a single example view of the face. For brevity, in the present paper we describe the application of the algorithm to the synthesis of a "frontal" view (i.e., defined in this paper as the novel view) from an example "rotated" view (i.e., defined in this paper as the view 24 ffi from frontal). It should be noted, however, that the algorithm is not at all restricted to a particular orientation of faces. The algorithm can be subdivided into three parts (for an overview see figure 3). ffl First, the texture and shape information in an image of a face are separated. ffl Second, two separate modules, one for texture and one for shape, compute the texture and shape representations of a given "ro- tated" view of a face (in terms of the appropriate view of the reference face). These modules are then used to compute the shape and texture estimates for the new "frontal" view of that face. ffl Finally the new texture and shape for a "frontal" view are combined and warped to the "frontal" image of the face. Separation of texture and shape in images of faces: The central part of the approach is a representation of face images that consists of a separate texture vector and 2D-shape vector, each one with components referring to the same feature points - in this case pixels. Assuming pixelwise correspondence to a reference face in the same pose, a given example image can be represented as fol- lows: its 2D-shape will be coded as the deformation field of n selected feature points - in the limit of each pixel - to the reference image. So the shape of a face image is represented by a vector that is by the y distance or displacement of each feature with respect to the corresponding feature in the reference face. The texture is coded as a difference map between the image intensities of the exemplar face and its corresponding intensities in the reference face. Thus, the mapping is defined by the correspondence field. Such a normalized texture can be written as a vector contains the image intensity differences i of the n pixels of the image. All images of the training set are mapped onto the reference face of the corresponding orientation. This is done separately for each rotated orientation. For real images of faces the pixelwise correspondences necessary for this mappings where computed automatically using a gradient based optical technique which was already used successfully previously on face images (Beymer et al., 1993; Vetter and Poggio, 1996). The technical details for this technique can be found in appendix B. Linear shape model of faces: The shape model of human faces used in the algorithm is based on the linear object class idea (the necessary and sufficient conditions are given in (Vetter and Poggio, is built on a training set of pairs of images of human faces. From each pair of im- ages, each consisting of a "rotated" and a "frontal" view of a face, the 2D-shape vectors s r for the "rotated" shape and s f for the "frontal"shape are computed. Consider the three-dimensional shape of a human head defined in terms of pointwise features. The 3D-shape of the head can be represented by a vector that contains the x; y; z-coordinates of its n feature points. Assume that S 2 ! 3n is the linear combination of q 3D shapes S i of other heads, such It is quite obvious that for any linear transformation R (e.g. rotation in 3D) Thus, if a 3D head shape can be represented as the weighted sum of the shapes of other heads, its rotated shape is a linear combination of the rotated shapes of the other heads with the same weights fi i . To apply this to the 2D face shapes computed from images, we have to consider the following. projection P from 3D to 2D with s under which the minimal number q of shape vectors necessary to represent i does not change, it allows the correct evaluation of the coefficients fi i from the images. Or in other words, the dimension of a three-dimensional linear shape class is not allowed to change under a projection P . Assuming such a projection, and that s r , a 2D shape of a given "ro- tated" view, can be represented by the "rotated" shapes of the example set s r i as then the "frontal" 2D-shape s f to a given s r can be computed without knowing S using fi i of equation (1) and the other s f given through the images in the training set with the following equation: In other words, a new 2D face shape can be computed without knowing its three-dimensional structure. It should be noted that no knowledge of correspondence between equation (1) and equation (2) is necessary (rows in a linear equation system can be exchanged freely). Texture model of faces: In contrast to the shape model, two different possibilities for generating a "frontal" texture given a "rotated" texture are de- scribed. The first method is again based on the linear object class approach and the second method uses a single three-dimensional head model to map the texture from the "rotated" texture onto the "frontal" texture. The linear object class approach for the texture vectors is equivalent to the method described earlier for the 2D-shape vectors. It is assumed that a "rotated" texture T r can be represented by the q "rotated" textures T r computed from the given example set as follows: It is assumed further that the new texture T f can be computed using ff i of equation (3) and the other given through the "frontal" images in the training set by the following equation: The three-dimensional head model: Whereas the linear texture approach is satisfactory for generating new "frontal" textures for regions not visible Correspondence of parts Correspondence of pixels Reference Faces Figure 2: A three-dimensional model of a human head was used to render the reference images (column the linear shape and texture model. The model defines corresponding parts in the two images (column B) and also establishes pixelwise correspondence between the two views (column C). Such a correspondence allows texture mapping from one view (C1) to the other (C2). in the "rotated" texture, it is not satisfactory for the regions visible in both views. The linear texture approach is hardly able to capture or represent features which are particular to an individual face (e.g. freckles, moles or any similar distinct aspect of facial texture). Such features ask for a direct mapping from the given "rotated" texture onto the new "frontal" texture. However, this requires pixelwise correspondence between the two views (see (Beymer et al., 1993)) . Since all textures are mapped onto the reference face, it is sufficient to solve the correspondence problem across the the viewpoint change for the reference face only. A three-dimensional model of an object intrinsically allows the exact computation of a correspondence field between images of the object from different viewpoints, because the three-dimensional coordinates of the whole object are given, occlusions are not problematic and hence the pixels visible in both images can be separated from the pixels which are only visible from one viewpoint. A single three-dimensional model of a human head is incorporated into the algorithm for three different processing steps. 1. The reference face images used for the formation of the linear texture and 2D-shape representations were rendered from the 3D- model under ambient illumination conditions (see figure 2A). 2. The 3D-model was manually divided into separate parts, the nose, the eye and mouth region and the rest of the model. Using the projections of these parts, the reference images for different orientations could be segmented into corresponding parts for which the linear texture and 2D-shape representation could be applied separately (see next paragraph on "The shape and texture models applied to parts" and also figure 2B). 3. The correspondence field across the two different orientations was computed for the two reference face images based on the given 3D- model. So the visible part of any texture, mapped onto the reference face in one orien- tation, can now be mapped onto the reference face in the second orientation (see figure 2C and 3). To synthesize a complete texture map on the "frontal" reference face for a new view, (i.e., the regions invisible in the exemplar view are lacking), the texture of the region visible in both views, which has been obtained through direct texture mapping across the viewpoint change, is merged with the texture obtained through the linear class approach (see figure 3). The blending technique used to merge the regions is described in detail in the appendix D. The shape and texture models applied to parts. The linear object class approach for 2D-shape and texture, as proposed in (Vetter and Poggio, 1996), can be improved through the 3D-model of the reference face. Since the linear object class approach CONSTRUCTED TEXTURE COMBINED TEXTURE LINEAR COMBINATION OF FRONTAL TEXTURES OUTPUT IMAGES REAL FACE l i INPUT IMAGE NORMALIZED TEXTURE MAPPED TEXTURE Input - S l T I i) INPUT SHAPE S Input CONSTUCTED SHAPE SF of Frontal View LINEAR COMBINATION OF FRONTAL SHAPES APPROXIMATION APPROXIMATION MAPPING INPUT IMAGE ONTO REFERENCE IMAGE Figure 3: Overview of the algorithm for synthesizing a new view from a single input image. After mapping the input image onto a reference face in the same orientation, texture and 2D-shape can be processed separately. The example based linear face model allows the computation of 2D-shape and texture of a new "frontal" view. Warping the new texture along the new deformation field (coding the shape) results in the new "frontal" views as output. In the lower row on the right the result purely based on the linear class approach applied to parts is shown, in the center the result with texture mapping from the "rotated" to the "frontal" view using a single generic 3D model of a human head. On the bottom left the real frontal view of the face is shown. did not assume correspondence between equations (1) and (2) or (3) and (4), shape and texture vectors had to be constructed for the complete face as a whole. On the other hand, modeling parts of a face (e.g. nose, mouth or eye region ) in independent separate linear classes is highly prefer- able, because it allows a much better utilization of the example image set and therefore gives a much more detailed representation of a face. A full set of coefficients for shape and texture representation is evaluated separately for each part instead of just one set for the entire face. To apply equations (1 - 4) to individual parts of a face, it is necessary to isolate the corresponding it areas in the "rotated' and ``frontal'' reference images. Such a separation requires the correspondence between the "rotated' and ``frontal'' reference image or equivalent between equations (1) and (2) of the shape representation and also between equations (3) and (4) for the texture. The 3D-model, however, used for generating the reference face images determines such a correspondence immediately (for example see figure 2B) and allows the separate application of the linear class approach to parts. To generate the final shape and texture vector for the whole face, this separation adds only a few complexities to the computational process . Shape and texture vectors obtained for the different parts must be merged, which requires the use of blending techniques to suppress visible border effects. The blending technique used to merge the regions is described in detail in appendix D. The algorithm was tested on 100 human faces. For each face, images were given in two orientations (24 ffi and 0 ffi ) with a resolution of 256-by-256 pixels and 8 bit (more details are given in appendix A). In a leave-one-out procedure, a new "frontal" view of a face was synthesized to a given "rotated" view (24 ffi ). In each case the remaining 99 pairs of face images were used to build the linear 2D- shape and texture model of faces. Figure 4 shows the results for six faces for three different implementations of the algorithm (center rows A,B,C). The left column shows the test image given to the algorithm. The true "frontal" view to each test face from the data base is shown in the right col- umn. The implementation used for generating the images in column A was identical to the method already described in (Vetter and Poggio, 1996), the linear object class approach was applied to the shape and texture vector as a whole, no partitioning of the reference face or texture mapping across the viewpoints was applied. The method used in was identical to A, except that the linear object class approach was applied separately to the different parts of a face. The three-dimensional head model was divided into four parts (see figure 2B) the eye, nose, mouth region, and the remaining part of the face. To segment the two reference images correctly, it was clearly necessary to render both of them from the same three-dimensional model of a head. Based on this segmentation, the texture and 2D-shape vectors for the different parts were separated and for each part a separate linear texture and 2D-shape model was ap- plied. The final image was rendered after merging the new shape and texture vectors of the parts. The images shown in column C are the result of a combination of the technique described in B and texture mapping across the viewpoint change. After mapping a given "rotated" face image onto the "rotated" reference image, this normalized texture can be mapped onto the "frontal" reference face since the correspondence between the two images of the reference face is given through the three-dimensional model. The part of the "frontal" texture not visible in the "rotated" view is substituted by the texture obtained by the linear texture model as described under B. The quality of the synthesized "frontal" views was tested in a simple simulated recognition experi- ment. For each synthetic image, the most similar frontal face image in the data base of 130 faces was computed. For the image comparison, two common similarity measures were used: a) the correlation coefficient, also known as direction cosine; and b) the Euclidean distance (L 2 ). Both measures were applied to the images in pixel representation without further processing. The recognition rate of the synthesized images (type A,B,C) was 100 % correct, both similarity measures independently evaluated the true "frontal" view to a given "rotated" view of a face as the most similar image. This result holds for all three different methods applied for the image syn- thesis. The similarity of the synthetic images to the real face image improved by applying the linear object class approach separately to the parts and improved again adding the correspondence between the two reference images to the method. This improvement is indicated in figure 5 where decreases where as the correlation coefficients increase for the different techniques. INPUT ROTATED SYNTHESIZED IMAGES Figure 4: Synthetic new frontal views (center columns) to a single given rotated (24 ffi ) image of a face (left column) are shown. The prior knowledge about faces was given through a training set of 99 pairs of images of different faces (not shown) in the two orientations. Column A shows the result based purely on the linear object class approach. Adding a single 3D-head model, the linear object class approach can be applied separately to the nose, mouth and eye region in a face (column B). The same 3D-model allows the texture mapping across the viewpoint change (column C). The frontal image of the real face is shown in the right column. Average Image Distance to Nearest Neighbor Real Face Images 4780.3 0.9589 Synthetic Images Type A 3131.9 0.9811 Synthetic Images Type B 3039.3 0.9822 Synthetic Images Type C 2995.0 0.9827 Figure 5: Comparing the different image synthesis techniques using Direction Cosines and L2-Norm as distance measures. First, for all real frontal face images the average distance to its nearest neighbor (an image of a different computed over an images test set of 130 frontal face images. Second, for all synthetic images (type A,B,C) the average value to its nearest neighbor was computed for both distance measures. For all synthetic images the real face image was found as nearest neighbor. Switching from technique A to B and from B to C the average values of Direction Cosines increase whereas the values of the L2-Norm decrease, indicating an improved image similarity. A crucial test for the synthesis of images is a direct comparison of real and synthetic images by human observers. In a two alternative forced choice subjects were asked to decide which of the two frontal face images matches a given rotated image (24 ffi ) best. One image was the "real" face the other a synthetic image generated applying the linear class method to the parts of the faces separately (method B). The first five images of the data set were used to familiarize the subjects with the task, whereas the performance was evaluated on the remaining 95 faces. Although there was no time limit for a response and all three images were shown simultaneously, there were only 6 faces classified correctly by all 10 subjects (see figure 6). In all other cases the synthetic image was at least by one subject classified as the true image and in one case the synthetic image was found to match the rotated image better as the real frontal image. In average each observer was 74% correct whereas the chance level was at 50%. The subjects responded in average after 12 seconds. The results demonstrate clearly an improvement in generating new synthetic images of a human face from only a single given example view, over techniques proposed previously (Beymer and Pog- gio, 1995; Vetter and Poggio, 1996). Here a single three-dimensional model of a human head was added to the linear class approach. Using this model the reference images could be segmented into corresponding parts and additionally any texture on the reference image could be mapped precisely across the view point change. The information used from the three-dimensional model is equivalent to the addition of a single correspondence field across the viewpoint change. This addition increased the similarity of the synthesized image to the image of the real face for the shape as well as for the texture. The improvement could be demonstrated in automated image comparison as well as in perceptual experiments with human observers. The results of the automated image comparison indicate the importance of the proposed face model for viewpoint independent face recognition sys- tems. Here the synthetic rotated images were compared with the real frontal face image. It should also be noted, that coefficients, which result from the decomposition of shape and texture into example shapes and textures, already give us a representation which is invariant under any 3D affine transformation, supposing of course the linear face model holds a good approximation of the target The difficulties experienced by human observers in distinguishing between the synthetic images and the real face images indicate, that a linear face model of 99 faces segmented into parts gives a good approximation of a new face, it also indicates possible applications of this method in computer graphics. Clearly, the linear model depends on the given example set, so in order to represent from a different race or a different age group, the model would clearly need examples of these, an effect well known in human perception (cf. e.g. (O'Toole et al., 1994)). The key step in the proposed technique is a dense correspondence field between images of faces seen from the same view point. The optical flow technique used for the examples shown worked well, however, for images obtained under less controlled conditions a more sophisticated method for finding the correspondence might be necessary. New Classification of Synthetic Versus Real Face Images Number of 6 17 22 Faces Figure For 95 different faces a rotated image (24 ffi ) and two frontal images were shown to human observers simultaneously. They had to decide which of the frontal images was the synthesized image (type B) and which one was the real image. The table shows the error rate for 10 observers and the related number of faces. In average each observer was correct in 74% of the trails (chance level was 50%) and the average response time was seconds. correspondence techniques based on active shape models (Cootes et al., 1995; Jones and Poggio, are more robust against local occlusions and larger distortions when applied to a known object class. There shape parameters are optimized actively to model the target image. Several open questions remain for a fully automated implementation. The separation of parts of an object to form separated subspaces could be done by computing the covariance between the pixels of the example images. However, for images at high resolution, this may need thousands of example images. The linear object class approach assumes that the orientation of an object in an image is known. The orientation of faces can be approximated computing the correlation of a new image to templates of faces in various orientations (Beymer, 1993). It is not clear jet how precisely the orientation should be estimated to yield satisfactory results. Appendix A Face Images. pairs of images of caucasian faces, showing a frontal view and a view taken 24 ffi from the frontal were available. The images were originally rendered for psychophysical experiments under ambient illumination conditions from a data base of three-dimensional human head models recorded with laser scanner (Cyberware TM ). All faces were without makeup, accessories, and facial hair. Ad- ditionally, the head hair was removed digitally (but with manual editing), via a vertical cut behind the ears. The resolution of the grey-level images was 256-by-256 pixels and 8 bit. Preprocessing: First the faces were segmented from the background and aligned roughly by automatically adjusting them to their two-dimensional centroid. The centroid was computed by evaluating separately the average of all x; y coordinates of the image pixels related to the face independent of their intensity value. A single three-dimensional model of a human head, recorded with a laser scanner (Cyberware TM ), was used to render the two reference images. B Computation of the To compute the 2D-shape vectors s r used in equations (1) and (2), which are the vectors of the spatial distances between corresponding points in the face images, the correspondence of these points has to be established first. That means we have to find for every pixel location in an image, e.g. a pixel located on the nose, the corresponding pixel location on the nose in the other image. This is in general a hard problem. However, since all face images compared are in the same orienta- tion, one can assume that the images are quite similar and occlusions are negligible. The simplified condition of a single view make it feasible to compare the images of the different faces with automatic techniques. Such algorithms are known from optical flow computation, in which points have to be tracked from one image to the other. We use a coarse-to-fine gradient-based gradient method (Bergen et al., 1992) and follow an implementation described in (Bergen and Hingo- rani, 1990). For every point x; y in an image I, the error term for y, with I x ; I y being the spatial image derivatives and ffi I the difference of intensity of the two compared images. The coarse-to-fine strategy refines the computed displacements when finer levels are processed. The final result of this computation (ffix; ffi y) is used as an approximation of the spatial displacement vector s in equation (1)and (2). The correspondence is computed towards the reference image from the example and test images. As a consequence, all vector fields have a common origin at the pixel locations of the reference image. C Linear shape and texture synthesis. First the optimal linear decomposition of a given shape vector in equation (1) and a given texture vector in equation (3) was computed. To compute the coefficients ff i (or similar fi i ) the "initial" vector T r of the new image is decomposed (in the sense of least square) to the q training image vectors given through the training images by min- imizing The numerical solution for ff i and fi i was obtained by an standard SVD-algorithm (Press and Flan- nery, 1992). The new shape and texture vectors for the "frontal" view were obtained through simple summation of the weighted "frontal" vectors (equations( 2) and (4)). D Blending of patches. Blending of patches is used at different steps in the proposed algorithm. It is applied for merging different regions of texture as well as for merging regions of correspondence fields which were computed separately for different parts of the face. Such a patch work might have little discontinuities at the borders between the different patches. It is known that human observers are very sensitive to such effects and the overall perception of the image might be dominated by these. For images Burt and Adelson (Burt and Adel- son, 1983; Burt and Adelson, 1985) proposed a multiresolution approach for merging images or components of images. First, each image patch is decomposed into bandpass filtered component im- ages. Secondly, this component images are merged separately for each band to form mosaic images by weighted averaging in the transition zone. Finally, these bandpass mosaic images are summed to obtain the desired composite image. This method was applied to merge the different patches for the texture construction as well as to combine the texture mapped across the viewpoint change with the missing part taken from the constructed one. Originally this merging method was only described for an application to images, however, the application to patches of correspondence fields eliminates visible discontinuities in the warped images. Taking a correspondence field as an image with a vector valued intensity, the merging technique was applied to the x and y components of the correspondence vectors separately. Synthesis of the New Image. The final step is image rendering. The new image can be generated combining the texture and shape vector generated in the previous steps. Since both are given in the coordinates of the reference image, for every pixel in the reference image the pixel intensity and coordinates to the new location are given. The new location generally does not coincide with the equally spaced grid of pixels of the destination image. The final pixel intensities of the new image are computed by linear interpola- tion, a commonly used solution of this problem known as forward warping (Wolberg, 1990). Acknowledgments I am grateful to H.H. B-ulthoff and T. Poggio for useful discussions and suggestions. Special thanks to Alice O'Toole for editing the manuscript and for her endurance in discussing the paper. I would like to thank Nikolaus Troje for providing the images and the 3D-model. --R Automatic creation of 3D facial mod- els Hierarchical motion-based frame rate conversion Face recognition under varying pose. Face recognition from one model view. Merging images through pattern decomposition. Active shape models - their training and application Parameterizing images for recognition and reconstruc- tion Robot vision. Motion and structure from orthographic projections. Synthetizing a color algorithm from examples. A novel approach to graphics. Extracting projective structure from single perspective views of 3D point sets. Analysis and synthesis of facial image sequences using physical and anatomical models. Digital actors for interactive television. Recognition by linear combinations of models. Symmetric 3D objects are an easy case for 2D object recog- nition Image synthesis from a single example image. The importance of symmetry and virtual views in three-dimensional object recogni- tion Image Warping. --TR --CTR Xiaoyang Tan , Songcan Chen , Zhi-Hua Zhou , Fuyan Zhang, Face recognition from a single image per person: A survey, Pattern Recognition, v.39 n.9, p.1725-1745, September, 2006 Philip L. Worthington, Reillumination-driven shape from shading, Computer Vision and Image Understanding, v.98 n.2, p.326-344, May 2005 A. Criminisi , A. Blake , C. Rother , J. Shotton , P. H. Torr, Efficient Dense Stereo with Occlusions for New View-Synthesis by Four-State Dynamic Programming, International Journal of Computer Vision, v.71 n.1, p.89-110, January 2007 Martin A. Giese , Tomaso Poggio, Morphable Models for the Analysis and Synthesis of Complex Motion Patterns, International Journal of Computer Vision, v.38 n.1, p.59-73, June 2000 Bernd Heisele , Thomas Serre , T. Poggio, A Component-based Framework for Face Detection and International Journal of Computer Vision, v.74 n.2, p.167-181, August 2007 Yongmin Li , Shaogang Gong , Heather Liddell, Constructing Facial Identity Surfaces for Recognition, International Journal of Computer Vision, v.53 n.1, p.71-92, June Athinodoros S. Georghiades , Peter N. Belhumeur , David J. Kriegman, From Few to Many: Illumination Cone Models for Face Recognition under Variable Lighting and Pose, IEEE Transactions on Pattern Analysis and Machine Intelligence, v.23 n.6, p.643-660, June 2001

face recognition;flexible templates;image synthesis;rotation invariance

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Quasi-Invariant Parameterisations and Matching of Curves in Images.

In this paper, we investigate quasi-invariance on a smooth manifold, and show that there exist quasi-invariant parameterisations which are not exactly invariant but approximately invariant under group transformations and do not require high order derivatives. The affine quasi-invariant parameterisation is investigated in more detail and exploited for defining general affine semi-local invariants from second order derivatives only. The new invariants are implemented and used for matching curve segments under general affine motions and extracting symmetry axes of objects with 3D bilateral symmetry.

Introduction The distortions of an image curve caused by the relative motion between the observer and the scene can be described by specific transformation groups (Mundy and Zisserman, 1992). For exam- ple, the corresponding pair of contour curves of a surface of revolution projected on to an image center can be described by a transformation of the Euclidean group as shown in Fig. 1 (a). If a planar object has bilateral symmetry viewed under perspective, the corresponding contour curves of the object in an image can be described by the special affine group (Kanade and Kender, 1983; Van Gool et al., 1995a) (see Fig. 1 (b)). The corresponding contour curves of a 3D bilateral symmetry such as a butterfly are related by a transformation of the general affine group (see Fig. 1 (c)). Since these corresponding curves are equivalent objects, they have the same invariants under the specific transformation group. Thus, invariants under these transformation groups are very important for object recognition and identification (Moons et al., 1995; Mundy and Zisserman, 1992; Pauwels et al., 1995; Rothwell et al., 1995; Gool et al., 1992; Weiss, 1993). Although geometric invariants have been studied extensively, existing invariants suffer from occlusion (Abu-Mostafa and Psaltis, 1984; Hu, 1962; Reiss, 1993; Taubin and Cooper, 1992), image noise (Cyganski et al., 1987; Weiss, 1988) and the requirement of point or line correspondences (Bar- rett and Payton, 1991; Rothwell et al., 1995; Zisserman et al., 1992). To cope with these problems, semi-local integral invariants were proposed (Bruckstein et al., 1993; Sato and Cipolla, 1996a; Sato and Cipolla, 1996b) recently. They showed that it is possible to define invariants semi- locally, by which the order of derivatives in invariants can be reduced from that of group curvatures to that of group arc-length, and hence the invariants are less sensitive to noise. As we have seen in these works, the invariant parameterisation guar- 118 Sato and Cipolla antees unique identification of corresponding intervals on image curves and enables us to define semi-local integral invariants even under partial occlusions. Although semi-local integral invariants reduce the order of derivatives required, it is known that the order of derivatives in group arc-length is still high in the general affine and projective cases (see table 1). In this paper, we introduce a quasi-invariant parameterisation and show how it enables us to use second order derivatives instead of fourth and fifth. The idea of quasi-invariant parameterisation is to approximate the group invariant arc-length by lower order derivatives. The new parameterisations are therefore less sensitive to noise, and are approximately invariant under a slightly restricted range of image distortions. The concept of quasi-invariants was originally proposed by Binford (Binford and Levitt, 1993), who showed that quasi-invariants enable a reduction in the number of corresponding points required for computing algebraic invariants. For example quasi-invariants require only four points for computing planar projective invariants (Binford and Levitt, 1993), while exact planar projective invariants require five points (Mundy and Zisser- man, 1992). It has also been shown that quasi- invariants exist even under the situation where the exact invariant does not exist (Binford and Levitt, 1993). In spite of its potential, the quasi-invariant has not previously been studied in de- tail. One reason for this is that the concept of quasiness is rather ambiguous and is difficult to formalise. Furthermore, the existing method is limited to the quasi-invariants based on point cor- GA GA SA Projective (8DOF) Special Affine Similarity (4DOF) (3DOF) Euclidean General Affine (6DOF) (d) (b) (a) (c) Fig. 1. Image distortion and transformation groups. A symmetric pair of contour curves (white curves) of (a) a surface of revolution, (b) planar bilateral symmetry and (c) 3D bilateral symmetry can be described by Euclidean, special affine (equi-affine) and general affine (proper affine) transformations under the weak perspective assumption. The image distortion caused by the relative motion between the observer and the scene can also be described by group transformations as shown in (d) and (e). Quasi-Invariant Parameterisations 119 ~ (a) (b) Fig. 2. Identifying interval of integration semilocally. (a) and (b) are images of a Japanese character extracted from the first and the second viewpoints. The interval of integration in these two images can be identified uniquely from invariant arc-length, w. For example, if w1 and e are a corresponding pair of points, then the interval [-1, 1] with respect to w corresponds to the interval [-1, 1] with respect to e w in the second image. Even though the curve is occluded partially (second image), the semi-local integral invariants can be defined on the remaining parts of the curve. respondences (Binford and Levitt, 1993), or the quasi-invariants under specific models (Zerroug and Nevatia, 1993; Zerroug and Nevatia, 1996). In this paper, we investigate quasi-invariance on smooth manifolds, and show that there exists a quasi-invariant parameterisation, that is a parameterisation approximately invariant under group transformations. Although the approximated values are no longer exact invariants, their changes are negligible for a restricted range of transforma- tions. Hence, the aim here is to find in parameterisations the best tradeoff between the error caused by the approximation and the error caused by image noise. Following the motivation, we investigate a measure of invariance which describes the difference from the exact invariant under group transforma- tions. To formalise a measure of invariance in differential formulae, we introduce the so called prolongation (Olver, 1986) of vector fields. We next define a quasi-invariant parameter as a function which minimises the difference from the exact invariant. A quasi invariant parameter under general affine transformations is then proposed. The proposed parameter is applied to semi-local integral invariants and exploited successfully for matching curves under general affine transformations in real image sequences. 2. Semi-Local Integral Invariants In this section, we review semi-local integral in- variants, and motivate the new parameterisation, quasi-invariant parameterisation. If the invariants are too local such as differential invariants (Cyganski et al., 1987; Weiss, 1988), they suffer from noise. If the invariants are too global such as moment (integral) invariants (Abu- Mostafa and Psaltis, 1984; Hu, 1962; Lie, 1927; Reiss, 1993; Taubin and Cooper, 1992), they suffer from occlusion and the requirement of corre- spondences. It has been shown recently (Sato and Cipolla, 1996a; Sato and Cipolla, 1996b) that it is possible to define integral invariants semi-locally, so that they do not suffer from occlusion, image noise and the requirement of correspondences. Consider a curve, C 2 R 2 , to be parameterised by t. It is also possible to parameterise the curve by invariant parameters, w, under specific transformation groups. These are called arc-length of the group. The important property of group arc-length in integral formulae is that it enables us to identify the corresponding interval of integration automatically. Consider a point, C(w 1 ), on a curve C to be transformed to a point, e on a curve e C by a group transformation as shown in Fig. 2. Since 1 e it is clear that if we take the same interval [01w; 1w] around C(w 1 ) and e these two intervals correspond to each other (see Fig. 2). That is, by integrating with respect to the group arc-length, w, the corresponding interval of integration of the original and the transformed curves can be uniquely identified. We now define semi-local integral invariants at point Z w1+1w w101w Fdw (1) where, F is any invariant function under the group. The choice of F provides various kinds of semi-local integral invariants (Sato and Cipolla, 1996b). If we choose the function F carefully, the integral formula (1) can be solved analytically, and the resulting invariants have simpler forms. For example, in the affine case, if we substitute then the integral formula is solved analytically, 120 Sato and Cipolla and the integral invariants can be described by: denotes the determinant of a matrix which consists of two column vectors, . The right hand side of (2) is actually the area made by two vectors, results have been proposed by Bruckstein (Bruck- stein et al., 1993) by a different approach. The important properties of semi-local integral invariants are as follows: 1. The limits of integration [01w;1w] in semi-local integral invariants are identified uniquely in original and transformed images from invariant parameterisations. Thus, we do not need to worry about the correspondence problem caused by a heuristic search of image features 2. Even though the curve is occluded partially as shown in Fig 2 (b), the semi-local integral invariants can be defined on the remaining parts of the curve. Thus, they do not suffer from the occlusion problem unlike classical moment invariants 3. In general, the lowest order differential invariant of a transformation group is a group curvature, and requires second, fourth, fifth and seventh order derivatives in Euclidean, special affine, general affine and projective cases (Guggenheimer, 1977; Olver et al., 1994). The semi-local integral invariants enable us to reduce the order of derivatives required from that of group curvatures to that of Table 1. Order of derivatives required for the group arc-length and curvature. In general derivatives more than the second order are sensitive to noise, and are not available from images. Thus, the general affine and projective arc-length as well as curvatures are not practical. group arc-length curvature Euclidean 1st 2nd special affine 2nd 4th general affine 4th 5th projective 5th 7th group arc-length. Since as shown in table 1, the order of derivatives of group arc-length is lower than that of group curvature, the semi-local integral invariants are more practical than differential invariants. From table 1 (Olver et al., 1994), it is clear that the semi-local integral invariants are useful under Euclidean and special affine cases, but they still require high order derivatives in general affine and projective cases. The distortion caused by a group transformation is often not so large. For exam- ple, the distortion caused by the relative motion between the observer and the scene is restricted because of the finite speed of the camera or object motions. In such cases, parameters approximated by lower order derivatives give us a good approximation of the exact invariant parameterisation. We call such a parameterisation a quasi-invariant parameterisation. In the following sections, we define the quasi-invariant parameterisation, and derive an affine quasi-invariant parameterisation. 3. Infinitesimal Quasi-Invariance We first derive the concept of infinitesimal quasi- invariance; that is quasi-invariance under infinitesimal group transformations. 3.1. Vector Fields of the Group Let G be a Lie group, that is a group which carries the structure of a smooth manifold in such a way that both the group operation (multiplica- tion) and the inversion are smooth maps (Olver, 1986). Transformation groups such as rotation, Euclidean, affine and projective groups are Lie groups. Consider an image point x to be transformed to an image point e by a group G: so that a function, I(x; y), with respect to x and y coordinates is transformed to ~ y) by h. Infinitesimally we can interpret this phenomenon by an action of a vector field, v: @ @x @ @y Quasi-Invariant Parameterisations 121 ~ ~ G Fig. 3. The vector field, v, and an integral curve, 0. The curve C is transformed to e C by a group transformation, so that the point P on the curve is transformed to e P. Locally the orbit of the point caused by a group transformation coincides with the integral curve, 0, of the vector field at the point, P. where, and j are functions of x and y, and provide various vector fields. Locally the orbit of the point, x, caused by the transformation, h, is described by an integral curve, 0, of the vector field, v, passing through the point (see Fig. 3). v is called an infinitesimal generator of the group ac- tion. The uniqueness of an ordinary differential equation guarantees the existence of such a unique integral curve in the vector field. Because of its linearity, any infinitesimal generator can be described by the summation of a finite number of independent vector fields, v i m), of the group as follows: is the ith independent vector field: @ @x @ @y are basis coefficients of @ @x and @ @y respectively, and are functions of x and y. These independent vector fields form a finite dimensional vector space called a Lie algebra (Olver, 1986). Locally any transformation of the group can be described by an integral of a finite number of independent vector fields, v i . The vector field described in (3) acts as a differential operator of the Lie derivative. 3.2. Exact Invariance We now state the condition of invariance of an arbitrary function I, which is well known in Lie Group theory. Let v be an infinitesimal generator of the group transformation. A real-valued function I is invariant under group transformations, if and only if the Lie derivative of I with respect to any infinitesimal generator, v, of the group, G, vanishes as follows (Olver, 1995): where denotes the Lie derivatives with respect to a vector field v. Since I is a scalar func- tion, the Lie derivative is the same as the directional derivative with respect to v. Thus, the condition of invariance (6) can be rewritten as follows: is the directional derivative with respect to v. 3.3. Infinitesimal Quasi-Invariance The idea of quasi-invariance is to approximate the exact invariant by a certain function I(x; y), which is not exactly invariant but nearly invariant. If the I , is not exactly invariant the equation no longer holds. We can however measure the difference from the exact invariant by using (7). By definition, the change in function I caused by the infinitesimal group transformation induced by a vector field, v, is described by the Lie derivative of I as follows: For measuring the invariance of a function irrespective of the choice of basis vectors, we consider an intrinsic vector field of the group, G. It is known (Sattinger and Weaver, 1986) that if the 122 Sato and Cipolla group is semi-simple (e.g. rotation group, special linear group), there exists a non-degenerate symmetric bilinear form called Killing form, K, of the Lie algebra as follows: where ad(v i ) denotes the adjoint representation 1 of v i , and tr denotes the trace. The Killing form provides the metric tensor, for the algebra: and the Casimir operator, C a , defined by the metric tensor is independent of the choice of the basis vectors: ij is the inverse of g ij . That is, the met- changes according to the choice of basis vectors, v i , so that C a is an invariant. Since g ij is symmetric, there exists a choice of basis vectors, m), by which g ij is diagonalised as follows: ae 61 if Such vector fields, v i are unique in the group, G, and thus intrinsic. By using the intrinsic vector fields in (8), we can measure the change in value of a function, I , which is intrinsic to the group, G. For measuring the quasi invariance of a function irrespective of the magnitude of the function, we consider the change in function, ffiI , normalised by the original function, I. We, thus, define a measure of infinitesimal quasi invariance, Q, of a function I by the squared sum of normalised changes in function caused by the intrinsic vector fields, I This is a measure of how invariant the function, I, is under the group transformation. If Q is small enough, we call I a quasi-invariant under infinitesimal group transformations. Unfortunately, if the group is not semi-simple (e.g. general affine group, general linear group), the Killing form is degenerate and we do not have such intrinsic vector fields. However, it is known that a non-semi-simple group is decomposed into a semi-simple group and a radical (Jacobson, 1962). Thus, in such cases, we choose a set of vector fields which correspond to the semi-simple group and the radical. 4. Quasi-Invariance on Smooth Manifolds In the last section, we introduced the concept of infinitesimal quasi-invariance, which is the quasi-invariance under infinitesimal group trans- formations, and derived a measure for the invariance of an approximated function. Unfortunately (11) is valid only for functions which do not include derivatives. In this section, we introduce an important concept known as the prolongation (Olver, 1986) of vector fields, and investigate quasi-invariance on smooth manifolds, so that it enables us to define quasi-invariants with a differential formula. 4.1. Prolongation of Vector Fields The prolongation is a method for investigating the differential world from a geometric point of view. Let a smooth curve C 2 R 2 be described by an independent variable x and a dependent variable y with a smooth function f as follows: The curve, C, is transformed to e C by a group transformation, h, induced by a vector field, v, as shown in Fig. 4. Consider a kth order prolonged space, whose coordinates are x, y and derivatives of y with respect to x up to kth order, so that the prolonged space is dimensional. The curves, C and e C, in 2D space are prolonged and described by space curves, C (k) and e in the dimensional prolonged space. The prolonged vector field, v (k) , is a vector field in k sion, which carries the prolonged curve, C (k) , to the prolonged curve, e explicitly as shown in Fig. 4. More precisely, the kth order prolonga- tion, v (k) , of a vector field, v, is defined so that it transforms the kth order derivatives, y (k) , of a into the corresponding kth order derivatives, e y (k) , of the transformed function geometrically. Quasi-Invariant Parameterisations 123 x y y x -0.4 x y x y y x -0.4 x y pr prolonged space ~ pr original image prolonged space transformed image Fig. 4. Prolongation of a vector field. The kth order prolonged vector field, v (k) transforms kth order derivatives of y into kth order derivatives of ey. That is the prolonged curve, C (k) , is transformed into the prolonged curve, e by the prolonged vector field, v (k) . This enables us to investigate derivatives of functions geometrically. pr (k) denotes kth order prolongation. This figure illustrates the first order prolongation m) be m independent vector fields induced by a group transformation, h. Since the prolongation is linear, the kth prolongation, of a general vector field, v, can be described by a sum of kth prolongations, v (k) i , of the independent vector fields, v i as follows: Consider a vector field (5) in 2D space again. Its first and second prolongations, v (1) , v (2) , are computed as follows (Olver, 1986): @y x v (2) @y xx where D x and D 2 x denote the first and the second total derivatives with respect to x, and y x , y xx , y xxx denote the first, second and the third derivatives of y with respect to x. Let F (x; function of x, y and derivatives of y with respect to x up to kth order, which is denoted by y (k) . Since the prolongation describes how the derivatives are going to change under group transformations, we can compute the change in function, ffiF , caused by the group transformation, h, as follows: where, v (k) is the kth order prolongation of the infinitesimal generator, v, of a transformation h. Note that we require only the same order of prolongation as that of the function, F . Since the prolongation describes how derivatives are going to change, it is important for evaluating the quasi- 124 Sato and Cipolla invariance of a differential formula as described in the next section. 4.2. Quasi-Invariance on Smooth Manifolds Let us consider the curve C in 2D space again. Suppose I(y (n) ) is a function on the curve containing the derivatives of y with respect to x up to the nth order, which we denote by y (n) . Since the nth order prolongation, v (n) , of the vector field v transforms nth order derivatives, y (n) , of the original curve to nth order derivatives, e y (n) , of the transformed curve, the change in function, caused by the infinitesimal group transformation induced by the ith independent vector described by: A quasi-invariant is a function whose variation caused by group transformations is relatively small compared with its original value. We thus define a measure of invariance, Q, on smooth curve, C, by the normalised squared sum of integrated along the curve, C, as follows: Z If I(y (n) ) is close to the exact invariant, then Q tends to zero. Thus, Q is a measure of how invariant the function, I(y (n) ), is under the group transformation. 5. Quasi-Invariant Parameterisation In the last section, we have derived quasi- invariance on smooth manifolds. We now apply the results and investigate the quasi-invariance of parameterisation under group transformations. A group arc-length, w, of a curve, C, is in general described by a group metric, g, and the independent variable, x, of the curve as follows: where, dw and dx are the differentials of w and x respectively. Suppose the metric, g, is described by the derivatives of y with respect to x up to kth order as follows: where y (k) denotes the kth order prolongation of y. The change of the differential, ffidw i , caused by the ith independent vector field, v i , is thus derived by computing the Lie derivative of dw with respect to the kth order prolongation of dx The change in dw normalised by dw itself is described as follows: ffid dw =g dx The measure of invariance of the parameter, w, is thus described by integrating the squared sum of ffid along the curve, C, as follows: Z where, dx If the parameter is close to the exact invariant parameter, Q tends to zero. Although there is no exact invariant parameter unless it has enough orders of derivatives, there still exists a parameter which minimises Q and requires only lower order derivatives. We call such a parameter a quasi-invariant parameter of the group, if it minimises (18) under the linear sum of the independent vector fields of the group and keeps Q small enough in a certain range of the group transformations. To find a function, g, which minimises (18) is in fact a variational problem (Gelfand and Formin, 1963) with the Lagrangian of L, which includes one independent variable, x, and two dependent variables, g and y (g is also dependent on y). In the next section, we derive a metric, g, which minimises the measure of invariance, Q, under general affine transformations by solving the variational problem. Quasi-Invariant Parameterisations 125 6. Affine Quasi-Invariant Parameterisa- tion In this section, we apply quasi-invariance to derive a quasi-invariant parameterisation under general affine transformations which requires only second order derivatives and is thus less sensitive to noise than the exact invariant parameter which requires fourth order derivatives. Suppose the quasi-invariant parameterisation, ' , under general affine transformation is of second order, so that the metric, ' g, of the parameter, , is made of derivatives up to the second: where, y x and y xx are the first and the second derivatives of y with respect to x. To find a quasi-invariant parameter is thus the same as finding a second order differential function, ' g(y x ; y xx ), which minimises the quasi-invariance, Q, under general affine transformations. Since the metric, ' g, is of second order, we require the second order prolongation of the vector fields to compute the quasi-invariance of the metric. 6.1. Prolongation of Affine Vector Fields A two dimensional general affine transformation is described by a 222 invertible matrix, A 2 GL(2), and a translational component, t 2 R 2 , and transforms e Since the differential form, d' , in (20) does not include x and y components, it is invariant under translations. Thus, we here simply consider the action of A 2 GL(2), which can be described by four independent vector fields, v i that is the divergence, curl, and the two components of deformation (Cipolla and Blake, 1992; Kanatani, 1990; Koenderink and van Doorn, @ @x @ @y @ @x @ @y @ @x @ @y @ @x @ @y Since the general linear group, GL(2), is not semi- simple, the Killing form (9) is degenerate and there is no unique choice of vector fields for the group (see section 3.3). It is however decomposed into the radical, which corresponds to the diver- gence, and the special linear group, SL(2), which is semi-simple and whose intrinsic vector fields coincide with (21). Thus, we use the vector fields in (21) for computing the quasi- invariance of differential forms under general affine transformations. From (12), (13), and (21), the second prolongations of these vector fields are computed by: v (2) @ @x @ @y @ @y xx v (2) @ @x @ @y @y x @ @y xx v (2) @ @x @ @y @ @y x @ @y xx v (2) @ @x @y @y x 03y x y xx @ @y xx These are the vector fields in four dimension, whose coordinates are x, y, y x and y xx , and the projection of these vector fields onto the x0y plane coincides with the original affine vector fields in two dimension. Since ' gdx is of second order, the prolonged vector fields, v (2) 3 and v (2) describe how the parameter, ' , is going to change under general affine transformations. 6.2. Affine Quasi-Invariant Parameterisation The measurement of the invariance, Q a , under a general affine transformation is derived by substituting the prolonged vector fields, v (2) 4 of (22) into (18): Z where, L is a function of y x , y xx , ' g and its derivatives as follows: v (2) dx 126 Sato and Cipollayg y x xx x * y Prolonged Space Image Fig. 5. Variation of ' g. The image curve C is transformed to C (n) in the prolonged space. The curve ' varies only on the surface 2 defined by C (n) . What we need to do is to find a curve ' g 3 on 2 which minimises Qa . Thus, there is no variation in yx and yxx . where, i denotes of ith vector field, v i , and ' g yx and ' g yxx are the first derivatives of ' g with respect to y x and y xx respectively. We now find a function, ' g, which minimises, Q a of (23). The necessary condition of Q a to have a minimum is that its first variation, ffiQ a vanishes: This is a variational problem (Gelfand and Olver, 1995) of one independent variable, x, and two dependent variables, y and ' g, and the integrand, L, of (23) is called the Lagrangian of the variational problem. It is known that (25) holds if and only if its Euler-Lagrange vanishes as follows (Olver, 1986): denotes the Euler operator. Since, in our case, one of the dependent variables, ' g, is a function of the derivatives, y x , y xx , of the other dependent variable, y, the Euler-Lagrange expression is different from the standard form of one independent and two dependent variables. We now investigate how this variational problem can be formalised. Suppose ' g changes to ' g+1'g, so that ' to ' g yx + 1'g yx and ' g yxx changes to ' respectively, where 1'g yx and 1'g yxx denote derivatives of 1'g with respect to y x and y xx . Then, the first variation of the function, Q a , caused by the change in ' g is described as follows: Z @'g @'g yx @'g yxx 1'g yxx Note that the variation occurs only on the surface 2 as shown in Fig. 5, and therefore the variation, ffiQ a does not include the change in y x and y xx . As shown in Appendix A, assuming that ' g has a form, y xx (where ff and fi are real values), we find that ffiQ a vanishes for any curve, if the following function, ' g, is chosen: We conclude that for any curve the following parameter ' is quasi-invariant under general affine transformations: The quasi-invariance, Q a , of an example curve computed by varying the power of y xx and (1+y 2 in (28) is shown in Fig. 6. We find that Q a takes a minimum when we choose ' shown in (28). By reformalising (28), we find that the parameter, ' , is described by the Euclidean arc-length, dv, and the Euclidean curvature, , as follows: Thus, d' , is in fact an exact invariant under ro- tation, and quasi-invariant under divergence and deformation. Note, it is known that the invariant parameter under similarity transformations is dv and that of special affine transformations is 1 The derived quasi-invariant parameter ' for general affine transformations is between these two as expected. We call ' the affine quasi-invariant parameter (arc-length). Since the new parameter requires only the second order derivatives, it is expected to be less sensitive to noise than the exact invariant parameter under general affine transformations Quasi-Invariant Parameterisations 127 betaalpha a a Fig. 6. Quasi-invariance of an artificial curve. The quasi-invariance, Qa , of an example curve is computed varying the power, ff and fl, in the parameter, dx. As we can see, Qa takes a minimum at This agrees with (28). 7. Quasi Affine Integral Invariants In this section, we apply the extracted affine quasi-invariant parameterisation to the semi-local integral invariants described in section 2, and derive quasi integral invariants under general affine transformations. defined in (29) is quasi-invariant, we can derive quasi integral invariants under general affine transformations by substituting ' in into w in (1). If we substitute F (' then we have the following quasi semi-local integral invariant: which is actually the area made by two vectors, points C(' 1 +1') and C(' 1 01') are identified by computing affine quasi-invariant arc-length, R d' . (30) is a relative quasi-invariant of weight one under general affine transformations as follows e can be computed from just second order derivatives, the derived invariants are much less sensitive to noise than differential invariants (i.e. affine curvature). This is shown in the experiment in section 9.2. 8. Validity of Quasi-Invariant Parameter- isation Up to now we have shown that there exists a quasi-invariant parameter under general affine transfor- mations, namely ' . In this section we investigate the systematic error of the quasi-invariant parameterisation, that is the difference from the exact invariance, and show under how wide range of transformations the quasi-parameterisation is valid. As we have seen in (29), the new parameterisation is an exact invariant under rotational motion. We thus investigate the systematic error 90 oo f y Fig. 7. Valid area of the affine quasi-invariant parame- terisation. The tilt and the slant motion of a surface is represented by a point on the sphere which is pointed by the normal to the oriented disk. The motion allowed for the affine quasi-invariant parameterisation is shown by the shaded area on the sphere, which is approximately less than in slant, and there is no preference in tilt angle. 128 Sato and Cipolla5030O O O O O (a) original curve (b) distorted curves Affine quasi-invariant arc-length Invariants O O (c) invariant signature Fig. 8. Systematic error in invariant signatures. The original curve on a fronto-parallel planar surface shown in (a) is distorted by the slant motion of the surface with tilt of 60 degrees (dashed lines) and slant of 30, 40, 50, and 60 degrees as shown in (b). The distortions of the curve caused by this slant motion can be modeled by general affine transformations. (c) shows invariant signatures made of quasi semi-local integral invariants (30) extracted from the curves in (b). If the slant motion of the plane is 40 degrees or more, the invariant signature suffers from large systematic error, while if the slant is less than 40 degrees, the proposed quasi-invariants are useful. caused by the remaining components of the affine transformation, that is the divergence and the deformation components. From (17), the systematic error of d' normalised by d' itself is computed to the first order: e =X a i v (2) dx where, a 1 , a 2 , a 3 and a 4 are the magnitude of di- vergence, curl, and deformation vector fields. Substituting (22) and (27) into (31), and since the curl component of the vector field does not cause any systematic error, we have: x 2y x x Since both (10y 2 x and 2yx in (32) vary only from 01 to 1, we have the following inequality: Thus, if ja 1 j 0:1, ja 3 j 0:1 and ja 4 j 0:1, then e 0:1, and the affine quasi-invariant parameterisation is valid. Fig. 7 shows the valid area of the affine quasi-invariant arc-length (parameteri- represented by the tilt and slant angles Quasi-Invariant Parameterisations 129 affine quasi-invariant arc-length integral invariants (a) Integral invariants (std 0.1) -0.020.020.06 affine arc-length differential invariants (b) Differential invariants (std 0.1) affine quasi-invariant arc-length integral invariants (c) Integral invariants (std 0.5) -0.020.020.06 affine arc-length differential invariants (d) Differential invariants (std 0.5) Fig. 9. Results of noise sensitivity analysis. The invariant signatures of an artificial curve are derived from the proposed invariants (semi-local invariants based on affine quasi invariant parameterisation) and the affine differential invariants (affine curvature), and are shown by thick lines in (a) and (b) respectively. The dots in (a) and (b) show signatures after adding random Gaussian noise of std 0.1 pixels, and the dots in (c) and (d) show signatures after adding random Gaussian noise of std 0.5 pixels. The thin lines show the uncertainty bounds of the signatures estimated by the linear perturbation method. The signatures from the proposed method are much more stable than those of differential invariants. which result in the systematic error, e, smaller than 0:1. In Fig. 7, we find that if the slant motion is smaller than 35 ffi , the systematic error is approximately less than 0.1, and the affine quasi-invariant parameterisation is valid. 9. Experiments 9.1. Systematic Error of Quasi Invariants In this section, we present the results of systematic error analysis of the quasi semi-local integral invariants, that is the semi-local integral invariants based on quasi-invariant parameterisation defined in (30), and show how large distortion is allowed for the quasi semi-local integral invariants. Fig. 8 (a) shows an image of a fronto-parallel planar surface with a curve. We slant the surface with tilt angle of 60 degrees and slant angle of 30, 40, 50 and 60 degrees as shown in Fig. 8 (b), and compute the invariant signatures of curves at each slant angle. Fig. 8 (c) shows invariant signatures of the curves computed from the quasi-invariant arc-length and the semi-local integral invariant (30). In this graph, we find that the invariant signature is distorted more under large slant motion as expected, and the proposed invariants are not valid if the slant motion is more than 40 degrees. Sato and Cipolla (a) viewpoint 1 (b) viewpoint 2 Affine quasi-invariant arc-length Invariants (c) invariant signature from viewpoint 1 Affine quasi-invariant arc-length Invariants (d) invariant signature from viewpoint 2 Fig. 10. Curve matching experiment. Images of natural leaves from the first and the second viewpoint are shown in (a) and (b). The white lines in these images show extracted contour curves. The quasi-invariant arc-length and semi-local integral invariants are computed from the curves in (a) and (b), and shown in (c) and (d) respectively. In this example, we chose It is clearly shown in these two signatures that the contour curve is partially occluded in (a). 9.2. Noise Sensitivity of Quasi Invariants We next compare the noise sensitivity of the proposed quasi semi-local integral invariants shown in (30) and the traditional differential invariants, i.e. affine curvature. The invariant signatures of an artificial curve have been computed from the proposed quasi- invariants and the affine curvature, and are shown by solid lines in Fig. 9 (a) and (b). The dots in (a) and (b) show the invariant signatures after adding random Gaussian noise of standard deviation of pixels to the position data of the curve, and the dots in (c) and (d) show those of standard deviation of 0.5 pixels. As we can see in these signa- tures, the proposed invariants are much less sensitive to noise than the differential invariants. This is simply because the proposed invariants require only second order derivatives while differential in- Quasi-Invariant Parameterisations 131 Affine quasi-invariant arc-length Invariants (a) invariant signatures (b) matched curves (c) matched curves Fig. 11. Results of curve matching experiment. The solid and dashed lines in (a) show the invariant signatures of the curves shown in Fig. 10 (a) and (b), which are shifted horizontally minimising the total difference between the two signatures. (b) and (c) show the corresponding curves extracted from the invariant signatures (a). variants require fourth order derivatives. The thin lines show the results of noise sensitivity analysis derived by the linear perturbation method. 9.3. Curve Matching Experiments Next we show preliminary results of curve matching experiments under relative motion between an observer and objects. The procedure of curve matching is as follows: 1. Cubic B-spline curves are fitted (Cham and Cipolla, 1996) to the Canny edge data (Canny, 1986) of each curve. This allows us to extract derivatives up to second order. 2. The quasi affine arc-length and the quasi affine semi-local integral invariants (30) with an arbitrary but constant 1' are computed at all points along a curve, and subsequently plotted as an invariant signature with quasi affine arc-length along the horizontal axis and the integral invariant along the vertical axis. The derived curve on the graph is an invariant signature up to a horizontal shift. We extract the invariant signatures of both the original and deformed curves. 3. To match curves we simply shift one invariant signature horizontally minimising the total difference between two signatures. 4. Corresponding points are derived by taking identical points on these two signatures. Even though a curve may be partially occluded or partially asymmetric, the corresponding points can be distinguished by the same procedure Fig. 10 (a) and (b) show the images of natural leaves taken from two different viewpoints. The white lines in these images show example contour 132 Sato and Cipolla Affine quasi-invariant arc-length Invariants (a) Affine quasi-invariant arc-length Invariants Affine quasi-invariant arc-length Invariants (c) Fig. 12. Comparison of signature. The invariant signatures in (a), (b) and (c) are computed from Fig. 10 (a) and (b) by using three different 1' , i.e. curves extracted from B-spline fitting. As we can see in these curves, because of the viewer motion, the curves are distorted and occluded partially. Since the leaf is nearly flat and the extent of the leaf is much less than the distance from the camera to the leaf, we can assume that the corresponding curves are related by a general affine transformation The computed invariant signatures of the original and the distorted curves are shown in Fig. 10 (c) and (d) respectively. One of these two signatures was shifted horizontally minimising the total difference between these two signatures (see Fig. 11 (a)). The corresponding points on the contour curves were extracted by taking identical points in these two signatures, and are shown in Fig. 11 (b) and (c). Note that the extracted corresponding curves are fairly accurate. In this experiment, we have chosen computing invariant signatures. For readers' reference, we in Fig. 12 compare the invariant signatures computed from three different 1' . 9.4. Extracting Symmetry Axes We next apply the quasi integral invariants for extracting the symmetry axes of three dimensional objects. Extracting symmetry (Brady and Asada, 1984; Friedberg, 1986; Giblin and Bras- sett, 1985; Gross and Boult, 1994; Van Gool et al., 1995a) of objects in images is very important for recognising objects (Mohan and Nevatia, 1992; Gool et al., 1995b), focusing attention (Re- isfeld et al., 1995) and controlling robots (Blake, Quasi-Invariant Parameterisations 133 Rotation l ~ ~ (a) (b) Fig. 13. Bilateral symmetry with rotation. The left and the right parts of an object with bilateral symmetry are rotated with respect to the symmetry axis, L in (a). The intersection point, O1 , of two tangent lines, l 1 and e l 1 , at corresponding points, P1 and e P1 , of a bilateral symmetry with rotation lies on the symmetry axes, L in (b). If we have N cross points, O the symmetry axis can be computed by fitting a line to these cross points, O1 , O2 , 1 1 1, ON . 1995) reliably. It is well known that the corresponding contour curves of a planar bilateral symmetry can be described by special affine transformations (Kanade and Kender, 1983; Van Gool et al., 1995a). In this section, we consider a class of symmetry which is described by a general affine transformation. Consider a planar object to have bilateral symmetry with an axis, L. Suppose the planar object can be separated into two planes at the axis, L, and is connected by a hinge so that two planes can rotate around this axis, L, as shown in Fig. 13 (a). The objects derived by rotating these two planes have a 3D bilateral symmetry. This class of symmetry is also common in artificial and natural objects such as butterflies and other flying in- sects. Since the distortion in images caused by a three dimensional motion of a planar object can be described by a general affine transformation, this class of symmetry can also be described by general affine transformations under the weak perspective assumption. Thus, the corresponding two curves of this symmetry have the same invariant signatures under general affine transformations. We must note the following properties: 1. The skewed symmetry proposed by Kanade (Kanade, 1981; Kanade and Kender, 1983) is a special case of this class of symmetry, where the rotational angle is equal to zero and the distortion can be described by a special affine transformation with determinant of 01. 2. Unlike the skewed symmetry of planar ob- jects, 3D bilateral symmetry takes both negative and positive determinant in its affine ma- trix. The positive means that the two planes are on the same side of projected symmetry axis, and the negative means that the planes are on opposite sides of the symmetry axis in the image. The extracted invariant signatures of corresponding curves of 3D bilateral symmetry are therefore either the same (i.e. positive determinant) or reflections of each other (i.e. negative determinant). 3. Unlike the skewed symmetry of planar ob- jects, the symmetry axis of 3D bilateral symmetry is no longer on the bisecting line of corresponding symmetric curves. Instead, the cross points of the tangent lines at corresponding points on the symmetric curves lie on the symmetry axis as shown in Fig. 13 (b). Thus the symmetry axis can be extracted by computing a line which best fits to the intersection points of corresponding tangent lines. We next show the results of extracting symmetry axes of 3D bilateral symmetry. Fig. 14 (a) shows an image of a butterfly (Small White) with a flower. Since the two wings of the but- 134 Sato and Cipolla (a) original image (b) contour curves -20002000 Affine quasi-invariant arc-length Invariants (c) invariant signature of the left wing -20002000 Affine quasi-invariant arc-length Invariants (d) invariant signature of the right wing Fig. 14. Extraction of axis of bilateral symmetry with rotation. (a) shows the original image of a butterfly (Small White), perched on a flower. (b) shows an example of contour curves extracted by fitting B-spline curves (Cham and Cipolla, 1996) to the edge data (Canny, 1986). The invariant signatures of these curves are computed from the quasi-invariant arc-length and semi-local integral invariants. (c) and (d) are the extracted signatures of the left and the right curves in (b). In this example, we chose terfly are not coplanar, the corresponding contour curves of the two wings are related by a general affine transformation as described above. Fig. 14 (b) shows example contour curves extracted from (a). Note that not all the points on the curves have correspondences because of the lack of edge data and the presence of spurious edges. Fig. 14 (c) and (d) shows the invariant signatures computed from the left and the right wings shown in Fig. 14 (b) respectively. (In this example, we chose computing semi-local integral invari- ants.) Since the signatures are invariant up to a shift, we have simply reflected and shifted one invariant signature horizontally minimising the total difference between two signatures (see Fig. 15 (a)). As shown in these signatures, semi-local invariants based on quasi-invariant parameterisation are Quasi-Invariant Parameterisations 135 -20002000 Affine quasi-invariant arc-length Invariants (a) invariant signatures (b) tangent lines (c) symmetry axis Fig. 15. Results of extracting symmetry axis of 3D bilateral symmetry. The solid and dashed lines in (a) show the invariant signatures of the curves shown in Fig. 14 (b), which are reflected and shifted horizontally minimising the total difference between two signatures. The black lines in (b) connect pairs of corresponding points extracted from the invariant signatures in (a). The white lines and the square dots show the tangent lines for the corresponding points and their cross points. The white line in (c) shows the symmetry axis of the butterfly extracted by fitting a line to the cross points. quite accurate and stable. Corresponding points are derived by taking the identical points on these two signatures, and shown in Fig. 15 (b) by connecting the corresponding points. Tangent lines at every corresponding pair of points are computed and displayed in Fig. 15 (b) by white lines. The cross points of every pair of tangent lines are extracted and shown in Fig. 15 (b) by square dots. The symmetry axis of the butterfly is extracted by fitting a line to the cross points of tangent lines and shown in Fig. 15 (c). Although the extracted contour curves include asymmetric parts as shown in Fig. 14 (b), the computed axis of symmetry agrees with the body of the butterfly quite well. Whereas purely global methods, e.g. moment based methods (Friedberg, 1986; Gross and Boult, 1994), would not work in such cases. These results show the power and usefulness of the proposed semi-local invariants and quasi-invariant parameterisation. 10. Discussion In this paper, we have shown that there exist quasi-invariant parameterisations which are not exactly invariant but approximately invariant under group transformations and do not require high order derivatives. The affine quasi-invariant parameterisation is derived and applied for matching of curves under the weak perspective assumption. Although the range of transformations is lim- ited, the proposed method is useful for many cases especially for curve matching under relative motion between a viewer and objects, since the movements of a camera and objects are, in general, 136 Sato and Cipolla limited. We now discuss the properties of the proposed parameterisation. 1. Noise Sensitivity Since quasi-invariant parameters enable us to reduce the order of derivatives required, they are much less sensitive to noise than exact invariant parameters. Thus using the quasi-invariant parameterisation is the same as finding the best tradeoff between the systematic error caused by the approximation and the error caused by the noise. The derived parameters are more feasible than traditional invariant parameters. 2. Singularity The general affine arc-length (Olver et al., 1994) suffers from a singularity problem. That is, the general affine arc-length goes to infinity at inflection points of curves, while the affine quasi-invariant parameterisation defined in (29) does not. This allows the new parameterisation to be more applicable in practice 3. Limitation of the Amount of Motion As we have seen in section 8, the proposed quasi-invariant parameter assumes the group motion to be limited to a small amount. In the affine case, this limitation is about a 1 0:1, a 3 0:1 and a 4 0:1 for the divergence and the deformation components (there is no limitation on the curl component, a 2 ). Since, in many computer vision applications, the distortion of the image is small due to the limited speed of the relative motion between a camera and the scene or the finite distance between two cameras in a stereo system, we believe the proposed parameterisation can be exploited in many applications. Appendix A In this section, we derive the affine quasi-invariant arc-length, ' . As we have seen in (26), ffiQ a is described as follows: Z @'g @'g yx @'g yxx 1'g yxx dx By computing the Lie derivatives of ' g with respect to v (2) 3 and v (2) 4 in (22), we have: v (2) @'g @y xx v (2) @y x @'g @y xx v (2) @'g @y x @'g @y xx v (2) @y x @'g @y xx Note, that @'g @x and @'g @y components vanish. This is because ' g does not include x and y components, and by definition the prolonged vector fields act only on the corresponding components given explicitly (e.g. @ @x acts only on x component and does not act on y, y x or other components). Substituting into (24), we find that the La- grangian, L, is computed from: @y x 04'gy xx (2+3y 2 @y xx @y xx @y x @y x @'g @y xx Consider a derivative: d dy x @'g yx 1'g dx dy x dy x @'g yx 1'g dx dy x @'g yx 1'g yx dx dy x @'g yx dy 2 x By integrating both sides of (A4) with respect to dy x , the second term of (A1) can be described by: Z @'g yx @L @'g yx 1'g dx dy x aZ d dy x @'g yx Z @'g yx 1'gd dx dy x Similarly the third term of (A1) can be described by: Z @'g yxx @L @'g yxx 1'g dx dy xx aZ d dy xx @'g yxx Z @'g yxx 1'gd dx dy xx Quasi-Invariant Parameterisations 137 Thus, the variation, ffiQ a , is computed from (A5) and (A6) as follows: Z where, @L @'g yx 1'g dx dy x a @L @'g yxx 1'g dx dy xx a @'gd dy x @'g yx@L @'g yx d dx dx dy xd dy xx @'g yxx@L @'g yxx d dx dx dy xx where, a and b are the limit of integration specified by the curve, C. From (A3), the derivatives of L in (A8) and (A9) can be computed by: @'g @y x @y xx 024y x y xx (1+y 2 @y x @'g @y xx x )'g @'g @y x x )'g @'g @y xx @'g yx @y xx @y x @'g yxx @y xx x )'g @y x Since 1'g in (A7) must be able to take any value, ffiQ a vanishes if and only if: The question is what sort of function, ' g, makes the condition (A13) hold. Here, we assume that ' g takes the following form: and investigate the unknown parameters ff and fi for (A13) to hold for arbitrary curves. Substituting into (A10), (A11), and (A12), we have: @'g @'g yx @'g yxx Since y x and y xx in (A15), (A16) and (A17) take arbitrary values, the condition (A13) holds for arbitrary curves if: Thus, ff and fi must be: ff =5 Substituting (A19) and (A20) into (A14), we find that the following form for ' g gives the extremal to Acknowledgements The authors acknowledge the support of the EP- SRC, grant GR/K84202. Notes 1. The adjoint representation, ad(v i ), provides a matrix representation of the algebra, whose (j; component is described by a structure constant C j ik (Sat- tinger and Weaver, 1986). --R Recognitive aspects of moment invariants. General methods for determining projective invariants in imagery. A symmetry theory of planar grasp. Smoothed local symmetries and their implementation. Invariant signatures for planar shape recognition under partial occlusion. A computational approach to edge de- tection Automated B-spline curve representation with MDL-based active contours Surface orientation and time to contact from image divergence and deformation. In Sandini An affine transformation invariant curvature function. Finding axes of skewed symme- try Calculus of Variations. Local symmetry of plane curves. Analyzing skewed sym- metries Differential Geometry. Visual pattern recognition by moment in- variants Lie algebras. Recovery of the three-dimensional shape of an object from a single view Mapping image properties into shape constraints: Skewed symme- try Geometry of binocular vision and a model for stereopsis. Gesammelte Abhandlungen Perceptual organization for scene segmentation and description. Pattern Analysis and Machine Intelligence Foundations of semi-differential invariants Geometric Invariance in Computer Vision. Applications of Lie Groups to Differential Equations. Differential invariant signatures and flows in computer vision. Recognition of planar shapes under affine distortion. Recognizing planar objects using invariant image features. Planar object recognition using projective shape representation. Affine integral invariants and matching of curves. Affine integral invariants for extracting symmetry axes. Lie groups and algebras with applications to physics Object recognition based on moment (or algebraic) invariants. The characterization and detection of skewed symmetry. International Journal of Robotics Research In Projective invariants of shapes. Geometric invariants and object recog- nition Recognizing general curved objects efficiently. --TR --CTR Tat-Jen Cham , Roberto Cipolla, Automated B-Spline Curve Representation Incorporating MDL and Error-Minimizing Control Point Insertion Strategies, IEEE Transactions on Pattern Analysis and Machine Intelligence, v.21 n.1, p.49-53, January 1999

bilateral symmetry;integral invariants;semi-local invariants;differential invariants;curve matching;quasi-invariant parameterisations

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Rational Filters for Passive Depth from Defocus.

A fundamental problem in depth from defocus is the measurement of relative defocus between images. The performance of previously proposed focus operators are inevitably sensitive to the frequency spectra of local scene textures. As a result, focus operators such as the Laplacian of Gaussian result in poor depth estimates. An alternative is to use large filter banks that densely sample the frequency space. Though this approach can result in better depth accuracy, it sacrifices the computational efficiency that depth from defocus offers over stereo and structure from motion. We propose a class of broadband operators that, when used together, provide invariance to scene texture and produce accurate and dense depth maps. Since the operators are broadband, a small number of them are sufficient for depth estimation of scenes with complex textural properties. In addition, a depth confidence measure is derived that can be computed from the outputs of the operators. This confidence measure permits further refinement of computed depth maps. Experiments are conducted on both synthetic and real scenes to evaluate the performance of the proposed operators. The depth detection gain error is less than irrespective of texture frequency. Depth accuracy is found to be 0.51.2% of the distance of the object from the imaging optics.

Introduction A pertinent problem in computational vision is the recovery of three-dimensional scene structure from two-dimensional images. Of all problems studied in vision, the above has, by far, attracted the most attention. This has resulted in a variety of sensors and algorithms [Jarvis-1983, Besl-1988] that can be broadly classified into two categories: active and passive. Active techniques produce relatively reliable depth maps, and have been applied to many industrial applications. However, when the environment cannot be controlled, as in the case of distant objects in outdoor scenes, active methods prove impractical. As a consequence, passive techniques are always desirable. Passive sensing methods, such as stereo and structure from motion, rely on algorithms that establish local correspondences between two or more images. From the resulting disparity estimates or motion vectors, the depths of points in the scene are computed. The process of determining correspondence is widely acknowledged as being computationally expensive. In addition, the above techniques suffer from the occlusion or missing part problems; it is not possible to compute depths of scene points that are visible in only one of the images. Alternative passive techniques are based on focus anal- ysis. Depth from focus uses a sequence of images taken by changing the focus setting of the imaging optics in small steps. For each pixel, the focus setting that maximizes image contrast is determined. This, in turn, can be used to compute the depth of the corresponding scene point [Horn-1968, Jarvis-1983, Krotkov-1987, Darrell and Wohn-1988, Nayar and Nakagawa-1994] . In contrast, depth from defocus uses only two images with different optical settings [Pentland-1987, Subbarao-1988, Ens and Lawrence-1991, Bove, Jr.-1993, Subbarao and Surya-1994, Nayar et al.-1995, Xiong and Shafer-1995]. The relative defocus in the two images can, in principle, be used to determine three-dimensional structure. The focus level in the two images can be varied by changing the focus setting of the lens, by moving the image sensor with respect to the lens, or by changing the aperture size. Depth from defocus is not confronted with the abovementioned missing part and correspondence problems. This makes it an attractive prospect for structure estimation. Despite these merits, at this point in time, fast, accurate, and dense depth from defocus has only been demonstrated using active illumination that constrains the dominant frequencies of the scene texture [Nayar et al.-1995, Watanabe et al.-1995] . Past investigations of passive depth from defocus indicate that it can prove computationally expensive to obtain a reliable depth map. This is because the frequency characteristics of scene textures are, to a large extent, unpredictable. Furthermore, the texture itself can vary dramatically over the image. Since the response of the defocus (blur) function varies with texture frequency, a single broadband filter that produces an aggregate estimate of defocus for an unknown texture cannot lead to accurate depth estimates. The obvious solution is to use an enormous bank of narrow-band filters and compute depth in a least-squares sense using all dominant frequencies of the texture [Xiong and Shafer-1995, Gokstorp-1994] . This requires one to forego computational efficiency. To worsen mat- ters, a depth map of high spatial resolution can be obtained only if all the filters in the bank have small kernel sizes. The uncertainty relation [Bracewell-1965] tells us that the frequency resolution of the filter bank reduces proportional to the inverse of the kernel size used. In short, one cannot design a filter with narrow enough response if the support area of the filter kernel is small. Xiong and Shafer [Xiong and Shafer-1995] proposed an attractive way to cope with this problem. They used moment filters to compensate for the frequency spectrum of the texture within the passband of each of the narrowband filters. This approach results in accurate depth estimates but requires the use of four additional filters for each of the tuned filters in the filter bank. This translates to five times as many convolutions as is needed for any typical filter bank method. Xiong and Shafer [Xiong and Shafer-1995] use convolutions in total, which makes their approach computationally expensive. Ens and Lawrence [Ens and Lawrence-1991] have proposed a method based on a spatial-domain analysis of two blurred images. They estimate the convolution matrix, which is convolved with one of the two images to produce the other image. The matrix corresponds to the relative blur between the two images. Once the matrix is computed, it can be mapped to depth estimates. This method produces accurate depth maps. How- ever, the iterative nature of the convolution matrix estimation makes it computationally expensive. Subbarao and Surya [Subbarao and Surya-1994] proposed the S-Transform and applied it to depth from defocus. They modeled the image as a third-order polynomial in spatial domain, and arrived at a simple and elegant expression [Subbarao and Surya- 1994]: are the far and near focused images, respectively. The blur circle diameters in images i 1 and i 2 are expressed by their second central moments oe 2 2 and oe 1 respectively. Since an additional relation between oe 2 and oe 1 can be obtained from the focus settings used for the two images, oe 2 and oe 1 can be solved for and mapped to a depth estimate. As we see no terms that depend on scene frequency in equation (1), this can be considered to be a sort of texture-frequency invariant depth from defocus method. It produces reasonable depth estimates for large planar surfaces in the scene. However, it does not yield depth maps with high spatial resolution that are needed when depth variations in the scene are significant. We argue that this requires a more detailed analysis of image formation as well as the design of novel filters based on frequency analysis. In this paper, we propose a small set of filters, or operators, for passive depth from defocus. These operators, when used in conjunction, yield invariance to texture frequency while computing depth. The underlying idea is to precisely model relative image blur in frequency domain and express this model as a rational function of two linear combinations of basis functions. This rational expression leads us to a texture- invariant set of operators. The outputs of the operators are used as coefficients in a depth recovery equation that is solved to get a depth estimate. The attractive feature of this approach is that it uses only a small number of broadband linear operators with small kernel supports. Consequently, depth maps are computed not only with high efficiency and accuracy but also with high spatial resolution. Since our operators are derived using a rational expression to model relative image blur, they are referred to as rational operators. Rational operators are general, in that, they can be derived for any blur model. The paper is structured as follows. First, the concept of a texture invariant operator set is described. Next, all the operations needed for depth from defocus are discussed, including the use of prefiltering and coefficient smoothing. An efficient algorithm for obtaining a confidence measure from the operator outputs is outlined. These confidence measures are effectively used for further refinement of computed depth maps. In our specific implementation of rational operators, we have used three basis functions to model the relative blurring function. This has resulted in a set of three rational operators with kernel sizes of 7\Theta7. This operator set has been used to compute depth maps for both synthetic scenes and real scenes. The experimental results are analyzed to quantify the performance of the proposed depth from defocus approach. Depth From Defocus 2.1 Principle Fundamental to depth from defocus is the relationship between focused and defocused images[Born and Wolf-1965] . Figure 1 shows the basic image formation geometry. All light rays that are radiated by object point P and pass the aperture A are refracted by the lens to converge at point Q on the image plane. The relationship between the object distance d, focal length of the lens f , and the image distance d i is given by the lens law:d f Each point on the object plane is projected onto a single point on the image plane, causing a clear or focused image i f to be formed. If, however, the sensor plane does not coincide with the image plane and is displaced from it, the energy received from P by the lens is distributed over a patch on the sensor plane. The result is a blurred image of It is clear that a single image does not include sufficient information for depth estimation, as two different scenes defocused to different degrees could produce identical images. A solution to the depth estimation problem is achieved by using two images, separated by a known physical distance 2e [Ens and Lawrence-1991, Subbarao and Surya-1994]. The distance fl of the image i 1 from the lens should also be known. Given the above described setting, the problem is reduced to analyzing the relative blurring of each scene point in the two images and computing the position of its focused image. A restriction here is that the images of all of the scene points must lie between the far- focused sensor plane i 1 and the near-focused sensor plane i 2 . For ease of description, we introduce the normalized depth ff, which equals \Gamma1 at i 1 and 1 at i 2 . Then, using ff)e in the lens law (2), we obtain the depth d of the scene point. R f a A a O Figure 1: Image formation and depth from defocus. The two images, i 1 and i 2 , include all the information required to recover scene structure between the focused planes in the scene corresponding to the two images. 2.2 Defocus Function Precise modeling of the defocus function is critical to accurate depth estimation. The defocus function is described in detail in previous works [Born and Wolf-1965, Horn- 1986]. In Figure ff)e is the distance between the focused image of a scene point and its defocused image formed on the sensor plane. The light energy radiated by the scene point and collected by the imaging optics is uniformly distributed on the sensor plane over a circular patch with a radius of (1 \Sigma ff)e a=d i 1 . This distribution, also called the pillbox, is the defocus function: is used for image used for image i 2 , and \Pi(r) is the rectangular function which takes the value 1 for jrj ! 1and 0 otherwise. F e is the effective F-number of the optics. In the optical system shown in Figure 1, F e equals d i =2a. In order to eliminate magnification differences between the near and far focused images, we have used telecentric optics, which is described in Appendix 7.1.1 and detailed in [Watanabe and Nayar-1995b]. In the telecentric case, F e equals f=2a 0 . In Fourier domain, the defocus function in (3) is: where, J 1 is the first-order Bessel function of the first kind, and u and v denote spatial frequency parameters in the x and y directions, respectively 2 . As is evident from the above expression, defocus serves as a low-pass filter. The bandwidth of the filter decreases as the radius of the blur circle increases, i.e. as the plane of focus gets farther from the sensor plane. Figure 2 illustrates this effect. Figure 2(a) shows the image i f (x; y) formed at the focused plane and its Fourier spectrum I f (u; v). When the sensor plane is displaced by a distance (1 \Gamma ff)e, the defocused image i 2 is the convolution of the focused image with the pillbox h 2 (x; y), as shown in Figure 2(b). The effect of defocus in spatial This geometric model is valid as far as the image is not exactly focused, in which case, a wave optics model is needed to describe the point spread function. Further, it is assumed that lens induced aberrations are small compared to the radius of the blur circle [ Born and Wolf-1965 ] . 2 In the past, most investigators have used the Gaussian model instead of the pillbox model for the blur function. This is mainly to facilitate mathematical manipulations; the Fourier transform of a Gaussian function is also a Gaussian which can be converted into a quadratic function by using the logarithm. As we will see, in our approach to depth from defocus, any form of blur function can be used. and frequency domains can be written as: I 2 (u; Since ff can vary from point to point in the image, strictly speaking, we have a space-variant system that cannot be expressed as a convolution. Therefore, equation (5) does not hold in a rigorous sense. However, if we assume that ff is constant in a small patch around each pixel, equation (5) remains valid within the small patch. Hereon, when we use the terms Fourier transform or spectrum, they are assumed to be those of a small image patch. For the assumption that ff variation in a patch is small to be valid, the patch itself must be small. In practice, to realize this requirement, one is forced to use broadband filters; the kernel size of a linear filter is inversely proportional to the bandwidth of the filter. Figure 2(c) is similar to (b) except that the sensor lies at the distance (1 from the focused plane to produce the defocused image i 1 . Again: I 1 (u; Note that in the spectrum plots we have used the polar coordinates (f r ; f ' ) for spatial frequencies, rather than Cartesian coordinates (u; v). This is because the defocus function is usually rotationally symmetric. This symmetry allows us to express the defocus spectrum using a single parameter, namely, the radial frequency f We see in Figure 2 that, since the image in (c) is defocused more than the one in (b), the low-pass response of H 1 (u; v) is greater than that of H 2 (u; v). 2.3 Depth from Two Images We now introduce the normalized ratio, M and P (u; Equivalently, in the spatial domain, we have the spectrum I f (u; v) of the focused image, which appears in equations (5) and (6), gets cancelled, the above normalized ratio is simply: (1- a) e /2Fe x y 1.22Fe /(1- a)e 4Fe /p (1- a) e (1+ a) e /2Fe x y 4Fe /p (1+ a) e if FT FT FT 1.22Fe /(1+ a)e (a) (b) (c)2 If Figure 2: The effect of blurring on the near and far focused images. (a) Focused image i f and its Fourier spectrum. (b) Pillbox defocus model h 2 and the Fourier spectrum I 2 of the blurred image. (c) Pillbox defocus model h 1 and the Fourier spectrum I 1 of the image for larger blurring. f is the radial frequency. Figure 3 shows the relationship between the normalized image ratio M=P and the normalized depth ff for several spatial frequencies. It is seen that M/P is a monotonic function of ff for large. As a rule of thumb, this frequency range equals the width of the main lobe of the defocus function H when it is maximally defocused, i.e. when the distance between the focused image i f and the sensor plane is 2e. From the zero-crossing of the defocus function in Figure 2, the highest frequency below which the normalized image ratio M/P is monotonic is found to be: For any given frequency within the above bound, since M/P is a monotonic function of ff, M/P can be unambiguously mapped to a depth estimate fi, as shown in Figure 3. a Figure 3: Relation between the normalized image ratio M=P and the defocus parameter ff. An upper frequency bound can be determined, below which, M=P is a monotonic function of the defocus parameter ff. For any given frequency within this bound, M=P can be unambiguously mapped to a depth estimate fi. Besides serving a critical role in our development, Figure 3 also gives us new way of viewing previous approaches to depth from defocus: If one can by some method determine the amplitudes, I 1 and I 2 , of the spectra of the two defocused images at a predefined radial frequency f 2 , a unique depth estimate can be obtained. This is the basic idea that most of the previous work is based on [Pentland-1987, Gokstorp-1994, Xiong and Shafer-1995], although the ratio used in the past is simply I 1 =I 2 rather than the normalized ratio M=P introduced here. Magnitudes of the two image spectra, at a predefined frequency, can be determined using linear operators (convolution). However, this is not a trivial problem. The image texture is unknown and can include unpredictable dominant frequencies and hence it is not possible to fix a priori the frequency of interest. This problem may be resolved by using a large bank of narrowband filters that densely samples the frequency space to estimate powers at a large number of individual frequencies. However, important trade-offs emerge while implementing narrowband linear operators [Gokstorp-1994, Xiong and Shafer-1995]. First, such an approach is clearly inefficient from a computational perspective. Furthermore, the uncertainty relation [Bracewell-1965] tells us that, when we apply frequency analysis to a small image area, the frequency resolution reduces proportional to the inverse of the area used. To obtain a dense depth map, one must estimate H 1 I and H 2 I using a very small area around each pixel. A narrow filter in spatial domain corresponds to a broadband filter in frequency domain. As a result, any operator output is inevitably an average of the local image spectrum over a band of frequencies. Since the response of the defocus function H depends on the local depth ff, and is not uniform within the pass-band of the operator, the output of the operator is, at best, an approximate focus measure and can result in large errors in depth. Given that all linear operators, however carefully designed, end up having a pass- band, it would be desirable to have a set of broadband operators that together provide focus measures that are invariant to texture. Further, if the operators are broadband, a small number of them could cover the entire frequency space and avoid the use of an extensive filter bank. The result would be efficient, robust, and high-resolution depth estimation. In the next section, we describe a method to accomplish this. 3 Rational Operator Set 3.1 Modeling Relative Defocus using a Rational Expression We have established the monotonic response of the normalized image ratio M=P to the normalized depth (or defocus) ff over all frequencies (see equation (7) and Figure 3). Our objective here is to model this relation in closed form. In doing so, we would like the model to be precise and yet lead us to a small number of linear operators for depth recovery. To this end, we model the function M=P by a rational expression of two linear combinations of basis functions: are the basis functions, G P i (u; v) and are the coefficients which are functions of frequency (u; v), and "(u; v; ff) is the residual error of the fit of the model to the function M=P . If the model is accurate, the residual error is negligible, and it becomes possible to use the model to map the normalized image ratio M=P to the normalized depth ff. The above expression can be rewritten as: Here, ff on the left hand side represents the actual depth of the scene point while fi on the right is the estimated depth. A difference between the two can arise only when the residual error is non-zero. If the normalized ratio on the left side is given to us for any frequency (u; v), we can obtain the depth estimate fi by solving equation (10). The above model for the normalized image ratio is general. In principle, any basis that captures the monotonicity and structure of the normalized ratio can be used. To be specific in our discussion, we use the basis we have chosen in our implementation. Since the response of M=P to ff is odd-symmetric and is almost linear for small radial frequencies f r (see Figure 3), we could model the response using three basis functions that are powers of fi: Then, equation (10) becomes 3 : The term including fi 3 can be seen as a small correction that compensates for the discrepancy of M=P from a linear model. From the previous section, we know that the blurring model completely determines M=P for any given depth ff and frequency (u; v). 3 We found that replacing b P2 (ff) by a tanh aff) gives us a slightly better fit when the defocus model is the pillbox function. Yet, to reduce the computational cost of solving equation (10) for depth fi, we have chosen this simple polynomial model. The above polynomial model, R(fi; u; v), can therefore be fit to the theoretical M=P in equation (7) by assuming fi to be ff. This gives us the unknown ratios G P1 =GM1 and G P2 =GM1 as functions of frequency (u; v). In the case of a rotationally symmetric blurring model, such as the pillbox function, these ratios reduce to functions of just the radial frequency f r . Now, if we fix any one of the coefficient functions, say, G P1 (u; v), all the other coefficients can be determined from the ratios 4 . Therefore, it is possible to determine all the coefficient functions that ensure that the above polynomial model accurately fits the normalized image ratio M=P given by equation (7). Figure 4 shows an example set (based on an arbitrary selection of G P1 (u; v)) of the coefficient functions, G P1 , G P2 and GM1 , for the case of the pillbox blur model. In the general form of the rational expression in equation (9), the coefficients of the rational expression can only be determined up to a multiplicative constant at each frequency. Therefore, we have: Here, -(u; v) is the unknown scaling function of all the coefficient functions and G P i (u; v) and GM i (u; v) represent the structures of the ratios obtained by fitting R(fi; u; v) to ff). The frequency response of the unknown scaling function -(u; v) is needed to determine all the coefficient functions without ambiguity. How this can be accomplished for the general rational expression will be described in section 4.1. We now examine how well the polynomial model fits the plots in Figure 3 of the normalized ratio M precisely, we are interested in knowing how well the model can used to estimate depth. To this end, for each frequency, we select a "true" depth value ff and find the corresponding ratio M=P using the analytical expression in (7). This ratio is then plugged into the polynomial model of (12) to calculate the depth estimate fi using the Newton-Raphson method. This process is repeated for all frequencies. Let us rewrite equation (12) as: As the third-order term can be considered to be a small correction, the following initial In practice, GP1 (u; v) cannot be selected arbitrarily. There are other restrictions that need to be considered. The exact selection procedure is discussed later in section 4.1. GP1(fr, fq) GP2(fr, fq) GM1(fr, fq) Figure 4: An example set of the coefficient functions obtained by fitting the polynomial model to the normalized image ratio M=P . Here, G P1 (u; v) was chosen and the remaining two functions determined from the fit. INVARIANCE TO fr Figure 5: Depth fi, estimated using the polynomial model in equation (12), is plotted as a function of spatial frequency for different values of actual depth ff. We see that the estimated depth equals the actual depth and is invariant to frequencies within the upper bound f r max given by equation (17). value can be provided to the Newton-Raphson method: Then, the solution after one iteration is: Figure 5 shows that the estimated depth fi is, for all practical purposes, equal to the actual depth, indicating that the polynomial model is indeed accurate. Further, the estimated depth is invariant (insensitive) to texture frequency as far as the radial frequency f r is below f r max . Above this frequency limit f r max , the response of M to ff, shown in Figure 3, becomes non-monotonic within the region an accurate depth estimate is not obtainable. In practice, any image can be convolved using a passband filter to ensure that all frequencies above f r max are removed. The rule of thumb used to determine f r max is given by equation (8). However, for the pillbox blur model, we have found via numerical simulation that f r max is in fact 1.2 times larger 5 than the limit given by equation (8). e This is a valuable side-effect of introducing the normalized image ratio M=P ; we can utilize 20% more frequency spectrum information than conventional methods which use the ratio I 1 =I 2 . 3.2 Rational Operator Set We have introduced a rational expression model for the normalized ratio M=P and shown that the solution of equation (10) gives us robust depth estimates for all frequencies within a permissible range. Thus far, this robustness was demonstrated for individual frequencies. In this section, we show how the rational model can be used to design a small set of broadband operators that can handle arbitrary textures. 5 This number can be increased from 1.2 to 1.3 if a larger number of Newton-Raphson iterations are used. However, depth results in this additional range are not numerically stable in the presence of noise since the response curves of M=P tend to flatten out. Hence, we use only one iteration. Taking cross-products in equation (10), we get: By integrating over the entire frequency space, we get: where: Here, we invoke the power theorem [Bracewell-1965]: where, F (u; v) and G(u; v) are the Fourier transforms of functions f(x; y) and g(x; y), respectively. Since we are conducting a spatial-frequency analysis, that is, we are analyzing the frequency content in a small area centered around each pixel, the right hand side of equation (20) is nothing but a convolution. This implies that cM i (ff) and c P i (ff) are actually functions of (x; y) and can be determined by convolutions as: are the inverse Fourier transforms of GM i (u; v) and In short, all the coefficients needed to compute depth using the polynomial in equation (19) can be determined by convolving the difference image m(x; y) and the summed image p(x; y) with linear operators that are spatial domain equivalents of the coefficient functions. We refer these as rational operators. The outputs of these operators at each pixel (x; y) are plugged into equation (19) to determine depth fi(x; y). As an example, if we use the model in equation (12), the depth recovery equation becomes: By substituting equation (22), we have: Again, the above rational operators are nothing but inverse Fourier transforms of the coefficient functions shown in Figure 4. We see that, though the operators are all broad-band (see Figure 4), the above recovery equation is independent of scene texture and provides an efficient means of computing precise depth estimates. 4 Implementation of Rational Operators The previous section described the theory underlying rational operators. In this section, we discuss various design and implementation issues that must be addressed to ensure that the rational operators produce accurate depth from defocus. In particular, we describe the design procedure used to optimize rational operator kernels, the estimation of a depth confidence measure, prefiltering of images prior to application of the rational operators, and the post-processing of the outputs of the operators. Since the rational expression model of equation (9) is too general, we focus on the simpler model of equation (12) which we used in our experiments. However, the procedures described here can be applied to other forms of the rational model. 4.1 Design of Rational Operators Kernels Since our rational operators are broadband linear filters, we can implement them with small convolution kernels. This is beneficial for two reasons; (a) low computational cost and (b) high spatial resolution. However, as we shall see, the design problem itself is not trivial. Note that, after deriving the operators, the functions G equation must have a ratio that equals the one obtained by fitting the polynomial model to the normalized image ratio. Any discrepancy in this ratio would naturally cause depth estimation errors. Fortunately, the base form function -(u; v) of equation (13) remains at our discretion and can be adjusted to minimize such discrepancies. This does not imply that -(u; v) will be selected arbitrarily, but rather that it will be given a convenient initial form that can be optimized. Clearly, the effect of discrepancies in the ratio would vary with frequency and hence depend on the textural properties of the scene. The design of the operator kernels is therefore done by minimizing an objective function that represents ratio errors over all frequencies. The relation between depth estimation error and ratio error is derived in appendix 7.2. We argue in the appendix that, for the depth error to be kept at a minimum, the ratio errors must satisfy the following oe GM1 (u; oe G P2 (u; oe GM1 (u; v) and oe G P2 (u; v) determine the weighting functions to be used in the minimization of errors in GM1 (u; v) and G P2 (u; v). Here, - is a constant and in the derivation of these expressions we have set -(u; v) equal to G P1 (u; v), i.e. G P1 (u; Therefore, from equation (13) we have -(u; G P2 (u; Now, we are in a position to formulate our objective function for operator design as follows: oe GM1 oe GM1 P2 (u; v) are the actual ratios of the designed discrete kernels, GM1 (u; v) and G P2 (u; v) are the ratios obtained in the previous section by fitting the polynomial model to the normalized image ratio, and oe GM1 o is a constant used to ensure that the minimization of - 2 does not produce the trivial result of zero-valued operators. M1 (0; 0) is the actural DC response of the designed discrete kernel g M1 , and GM1 (0; is its initial value. In the above summation, the discrete frequency samples should be sufficiently dense. When the kernel size is n \Theta n, the frequency samples should be at least 2n \Theta 2n in order to avoid the Gibbs phenomenon [Oppenheim and Schafer-1989]. In our optimization, we use 32 \Theta 32 sample points for 7 \Theta 7 kernels. Since - 2 is non-linear, its minimization is done using the Levenberg-Marquardt algorithm [Press et al.-1992] . We still need to define P (u; v; ff) in equation (25), which is dependent on the unknown texture of the image. However, since P (u; v; ff) is only used to fix the weighting functions in equation (25), a rough approximation suffices. To this end, we assume the distribution of the image spectrum to be: In our optimization we have used which corresponds to Brownian motion 6 . Though P (u; v; ff) changes with ff, we can use the approximation P (u; v; 6 If we denote fractal dimension [ Peitgen and Saupe-1988 ] by D h , in the two dimensional case the The last issue concerns the base form function -(u; in equation (25). An initial selection can be made for this function that will be refined by the optimization of - 2 . As GM1 (u; must be 0 in order to realize GM1 (u; using a finite kernel. Also, G P1 (u; v) must be smooth (without rapid fluctuations) to obtain rational operators with small kernels. In our implementation, we have imposed rotational symmetry as an added constraint and used the Laplacian of Gaussian to initialize G P1 (u; v): G P1 (f r f peak f peak where, f peak is the radial frequency at which G P1 is maximum. This frequency is set to 0:4f Nyquist in our optimization. Once again, the above function is only used for initialization and is further refined by the optimization of - 2 . An example set of discrete rational operators obtained from the optimization of - 2 will be presented shortly. 4.2 Prefiltering We now discuss prefiltering that needs to be applied to the input images i 1 (x; y) and y). The purpose is to remove the DC component and very high frequencies before applying the rational filters. The DC component is harmful because a small change in the illumination, between the two images, can cause an unanticipated bias in the image m(x; y). Such a bias would propagate errors to the coefficient image cM1 (x; y) since the GM1 operator applied to m(x; y) is essentially a low-pass filter. This, in turn, would cause depth errors. At the other end of the spectrum, radial frequencies greater than f r max (see equation (17)) are also harmful as they violate the monotonicity property of M=P , which is needed for rational operators to work. Therefore, such high frequencies must also be removed. Although it is possible to embed the desired prefilter within the rational filters (given that prefiltering can be done using linear operators), we have chosen to use a separate prefilter for the following reason. Since the prefilter attempts to cut low and high frequencies, it tends to have a large kernel. Embedding such a prefilter in the rational operators would require the operators also to have large kernels, thus, resulting in low spatial resolution as well as unnecessary additional computations. holds true. D corresponds to the case of extreme fractal, D 1:5 corresponds to Brownian motion and D corresponds to a smooth image. Finally, corresponds to white noise (completely random image). 7 In equation (12), M=P is zero when j(u; v)j ! 0. Since ff can be non-zero, 1=GM1 (u; must be zero for equation (12) to be valid. As with the rational operators, the design of the prefilter can be posed as the optimization of an objective function. Let us define the desirable frequency response of the prefilter as f(u; v). For reasons stated earlier, this frequency response must cut both the DC component and high frequencies. In addition, the frequency response should be smooth and rotationally symmetric to ensure a small kernel size. A function with these desired properties is again the Laplacian of Gaussian given by the right hand side of equation (28), but using f We define the objective function as: oe pass oe stop is the frequency response of the designed prefilter kernel. oe pass and oe stop represent the weights assigned to the passband and the stopband regions of the prefilter, respectively. The stopband is . The Levenberg-Marquardt algorithm [Press et al.-1992] is used to determine the prefilter kernel that 4.3 An Example Set of Discrete Rational Filters Figures 6 and 7 show the kernels and their frequency responses for the rational operators and the prefilter, derived with kernel size set to 7\Theta7 and e=F pixels. In order to make the operators uniformly sensitive to textures in all directions, we imposed the constraint that the kernels must be symmetric with respect to the x and y axes as well as the lines These constraints reduce the number of degrees of freedom (DOF) in the kernel design problem. In the case of a 7\Theta7 kernel, the DOF is reduced to 10. This further reduces to 6 for a 6\Theta6 or a 5\Theta5 kernel. This DOF of 6 is too small to design operators with the desired frequency responses. Therefore, the smallest kernel size was chosen to be k 7. Note that the passband response of the prefilter in Figure 7 can be further refined if its kernel size is increased. The final design issue pertains to the maximum frequency f r max . Since the discrete Fourier transform of a kernel of size k s has the minimum discrete frequency period of 1=k s , it is difficult to obtain precisely any response in the frequency region below 1=k s . Further, the spectrum in this region is going to be suppressed by the prefilter as it is close to the DC component. Therefore, the maximum frequency f r max must be well above 1=k s . We express this condition as f r ks . Using equation (17), we obtain: 2e This condition can be interpreted as follows: The maximum blur circle diameter 2e=F e must be smaller than 73% of the kernel size k s . This is also intuitively reasonable as the kernel should be larger than the blur circle as it seeks to measure blur 8 . 4.4 Coefficient Image Smoothing By applying the prefilter and the rational operators in Figure 6 to the images m(x; y) and p(x; y), we obtain coefficients that can plugged into equation (23) to compute depth fi. However, a problem can arise in solving for depth. If c P1 (x; in equation (24) is close to zero, the depth estimate becomes unstable as is evident from the solution step in equation (15). Since the frequency response of g P1 (x; y) cuts the DC component (Figure 5 (a)), zero-crossings are usually common in the coefficient image c P1 (x; y; ff). It is also obvious that, for image areas with weak texture 9 , c P1 (x; approaches zero. To solve this problem, we apply a smoothing operator to the coefficient image. This enables us to avoid unstable depth estimates at zero-crossings in the coefficient image, which otherwise must be removed by some ad hoc post-filtering. To optimize this smoothing operation, so as to minimize depth errors, we need an analytic model of depth error. Using the depth recovery equation (23), we get: Here, we have dropped the parameter (x; y) for brevity. Solving for dfi, we get: As c P2 is only a small correction factor, the following approximation can be made: c P1 We denote the standard deviations (errors) of cM1 , c P1 and c P2 by oe cM1 , oe cM1 and oe cM1 , respectively. To simplify matters, it is assumed that the errors are independent of each other. Then, we get [Hoel-1971]: oe fi 8 Since the above conditions related to kernel size are rough, we suggest that the linearity of depth estimation be checked (using synthetic images) to find the best kernel size k s . Such an evaluation is reported in the experimental section. 9 Weak texture is equivalent to low spectrum power in the high frequency region. \Gamma0:00133 0:0453 0:1799 0:297 0:1799 0:0453 \Gamma0:00133 0:0453 0:4009 0:8685 1:093 0:8685 0:4009 0:0453 0:1799 0:8685 2:957 4:077 2:957 0:8685 0:1799 0:297 1:093 4:077 6:005 4:077 1:093 0:297 0:1799 0:8685 2:957 4:077 2:957 0:8685 0:1799 0:0453 0:4009 0:8685 1:093 0:8685 0:4009 0:0453 \Gamma0:00133 0:0453 0:1799 0:297 0:1799 0:0453 \Gamma0:00133C C C C C C C C C C C C A \Gamma0:03983 \Gamma0:09189 \Gamma0:198 \Gamma0:259 \Gamma0:198 \Gamma0:09189 \Gamma0:03983 \Gamma0:09189 \Gamma0:3276 \Gamma0:4702 \Gamma0:4256 \Gamma0:4702 \Gamma0:3276 \Gamma0:09189 \Gamma0:259 \Gamma0:4256 1:393 3:385 1:393 \Gamma0:4256 \Gamma0:259 \Gamma0:09189 \Gamma0:3276 \Gamma0:4702 \Gamma0:4256 \Gamma0:4702 \Gamma0:3276 \Gamma0:09189 0:05685 \Gamma0:02031 \Gamma0:06835 \Gamma0:06135 \Gamma0:06835 \Gamma0:02031 0:05685 \Gamma0:02031 \Gamma0:06831 0:05922 0:1454 0:05922 \Gamma0:06831 \Gamma0:02031 \Gamma0:06835 0:05922 0:1762 \Gamma0:01998 0:1762 0:05922 \Gamma0:06835 \Gamma0:06135 0:1454 \Gamma0:01998 \Gamma0:698 \Gamma0:01998 0:1454 \Gamma0:06135 \Gamma0:06835 0:05922 0:1762 \Gamma0:01998 0:1762 0:05922 \Gamma0:06835 \Gamma0:02031 \Gamma0:06831 0:05922 0:1454 0:05922 \Gamma0:06831 \Gamma0:02031 0:05685 \Gamma0:02031 \Gamma0:06835 \Gamma0:06135 \Gamma0:06835 \Gamma0:02031 0:05685C C C C C C C C C C C C A pref ilter \Gamma0:143 \Gamma0:1986 \Gamma0:1056 \Gamma0:07133 \Gamma0:1056 \Gamma0:1986 \Gamma0:143 \Gamma0:1986 \Gamma0:1927 0:01795 0:07296 0:01795 \Gamma0:1927 \Gamma0:1986 \Gamma0:1056 0:01795 0:2843 0:4601 0:2843 0:01795 \Gamma0:1056 \Gamma0:07133 0:07296 0:4601 0:6449 0:4601 0:07296 \Gamma0:07133 \Gamma0:1056 0:01795 0:2843 0:4601 0:2843 0:01795 \Gamma0:1056 \Gamma0:1986 \Gamma0:1927 0:01795 0:07296 0:01795 \Gamma0:1927 \Gamma0:1986 \Gamma0:143 \Gamma0:1986 \Gamma0:1056 \Gamma0:07133 \Gamma0:1056 \Gamma0:1986 \Gamma0:143C C C C C C C C C C C C A Figure Rational operator kernels derived using kernel size of 7\Theta7 and e=F pixels. Regardless of scene texture, passive depth from defocus can be accomplished using this small operator set. u0.010.03(a) g M1 (b) g P1 u0.0050.015(c) g P2 (d) prefilter Figure 7: Frequency responses of the rational operators shown in Figure 6. This expression is useful as it gives us an estimate of depth error. The inverse of this estimate, 1=oe fi 2 , can be viewed as a depth confidence measure and be used to combine adjacent depth estimates in a maximum likelihood sense to obtain more accurate depth estimates. Also, when one wishes to apply depth from defocus at different scales using a pyramid framework [Jolion and Rosenfeld-1994, Burt and Adelson-1983, Darrell and Wohn-1988, Gokstorp-1994] , the above confidence measure can be used to combine depth values at different levels of the pyramid. In equation (34), oe c M1 , oe c P1 and oe c P2 are constants because they are defined by the readout noise of the image sensor used and the frequency responses of the rational operators. On the other hand, fi can be assumed to be locally constant, since depth can be expected to vary smoothly at most points in the image. These facts lead us to: With the above error model in place, we can develop a method for coefficient image smoothing. If we multiply equation (23) by c P1 (x; y; ff), and sum up depth values in the neighborhood R of each pixel, we get: c P1 (x; c P1 (x; where, fi a is the depth estimate after coefficient smoothing. Since the last terms in equations (36) and (23) are small corrections, fi a can be approximated by: fi a ' c P1 c P1 (x; c P1 c P1 From equation (35), we see that fi a is the weighted average of raw depth estimates fi in the neighborhood R, where the weights are 1=oe fi 2 (x; From statistics [Hoel-1971] we know that the optimal weighted average of independent variables whose variances are oe i 2 , is obtained by weighting the X i with 1=oe i 2 . Therefore, the above weighted average of depth estimates can be viewed as optimal. The variance oe a 2 of the resulting depth estimate fi a is given by:oe a 2 Hence, the coefficient smoothing of equation (36) is optimal, in that, it minimizes 10 the error in estimated depth fi a . In addition, the resulting smoothed coefficient c P1 (x; is proportional to the inverse of the variance of fi a , i.e. 1=oe fi a 2 , which is clear from equations (35), (37) and (38). Therefore, the smoothed coefficient c P1 (x; y; ff) can be used as a confidence measure to post-process computed depth maps. Another method to cope with zero-crossings in the cP1 coefficient image is based on the Hilbert This approach is detailed in [ Watanabe and Nayar-1995a PREFILTER PREFILTER FAR FOCUSED IMAGE NEAR FOCUSED IMAGE RATIONAL OPERATOR COEFFICIENT SMOOTHING DEPTH COMPUTATION Figure 8: The flow of the depth from defocus algorithm. Using Datacube's MV200 pipeline processor, the entire algorithm can be executed in as little as 0.16 msec to obtain a 512\Theta480 depth map. 4.5 Algorithm Figure 8 illustrates the flow of the depth from defocus algorithm we have implemented. The far and near focused images are first added and subtracted to produce p(x; y) and m(x; y), respectively. Then they are convolved with the prefilter and subsequently with the three rational operators. The resulting coefficient images are then smoothed by local averaging. The final step is the computation of depth from the coefficients using a single iteration of the Newton-Raphson method using equations (15) and (16). Alternatively, depth computation can be achieved using a precomputed two-dimensional look-up table. The look-up table is configured to take c 0 inputs and provides depth fi(x; y) as output. In summary, a depth map is generated with as few as 5 two-dimensional convolutions, simple smoothing of the coefficient images, and a straightforward depth computation step. The above operations can be executed efficiently using a pipelined image processor. If one uses Datacube's MV200 pipeline processor, all the computations can be realized using as few as 10 pipelines. The entire depth from defocus algorithm can then be executed in 0.16 msec for an image size of 512\Theta480. The efficiency of the algorithm, which comes from the use of the rational operator set, is far superior to any existing depth from defocus algorithm that attempts to compute accurate depth estimates [Xiong and Shafer-1993, Gokstorp-1994] . 5 Experiments 5.1 Experiments with Synthetic Images We first illustrate the linearity of depth estimation and its invariance to texture frequency using synthetic images. The synthetic images shown in Figure 9 correspond to a planar surface that is inclined away from the sensor such that its normalized depth value is 0 at the top and 255 at the bottom. The plane includes 10 vertical strips with different textural properties. The left 7 strips have textures with narrow power spectra whose central frequencies are 0.015, 0.03, 0.08, 0.13, 0.18, 0.25 and 0.35, from left to right. The th strip is white noise. The next two strips are fractals with dimensions of 3 and 2.5, respectively [Peitgen and Saupe-1988] . The near and far focused images were generated using the pillbox blur model. The defocus condition used was e=F pixels. In all our experiments, the digital images used are of size 640\Theta480. The depth map estimated using the 7\Theta7 rational operators and 5\Theta5 coefficient smoothing is shown as a gray- coded image in Figure 9(c) and a wireframe in Figure 9(d). As is evident, the proposed algorithm produces high accuracy despite the significant texture variations between the vertical strips. Figure summarizes quantitative results obtained from the above experiment. The figure includes plots of (a) the gradient of the estimated depth map, (b) RMS (root mean square) error (oe) in computed depth, and (c) the averaged confidence value. Each point (square) in the plots corresponds to one of the strips in the image, and is numbered from left to right (see numbers next to the squares). Note that the gradient of the estimated depth map is nothing but the depth detection gain. Figure 10(a) shows that the gain is invariant except for the left three strips. The slight gain error in the left three strips is because the ratio GM1 (u; v)=G P1 (u; v) is high for low frequencies. As a result, (a) far-focused image (b) near-focused image (c) gray-coded depth map (d) wireframe of depth map Figure 9: Depth from defocus applied to synthetic images of an the inclined plane is accurately recovered despite the significant texture variations. G P1 (u; v) is small in the low frequency region and a small error in G P1 (u; v) causes a large error in the ratio. The low values of G P1 (u; v) for low-frequency textures is reflected by the extremely low confidence values for the corresponding strips. However, as such low frequencies are cut by the prefilter, depth errors are suppressed if there exist other frequency components. When one wants to utilize low frequencies, a pyramid [Jolion and Rosenfeld-1994, Darrell and Wohn-1988] can be constructed and the rational filters can applied to each level of the pyramid. Depth maps computed at different levels of the pyramid can be combined in a maximum-likelihood sense using confidence measures which are easily computed along with the coefficient image using equation (34). Figure 10(b) and (c) show a rough agreement between the confidence measure plot and the function 1=oe 2 .0.2center freq.1.011.031.05gain 1/f 1/f 1/f (a)0.2center freq. (c) 1/f 1/f 1/f (b)0.40.8confidence 1/f 1/f 1/f 1.5freq. distribution confidence4080120246 Figure 10: Analysis of depth errors for the textured inclined plane shown in Figure 9. Each point (square) in the plots corresponds to a single texture strip on the inclined plane (numbered 1-10.) (a) The gradient of the computed depth map which corresponds to the depth detection gain. The invariance of depth estimation to image texture is evident. (b) The RMS error (oe) in computed depth. (c) The depth confidence value which is seen to be in rough agreement with 1=oe 2 . In Figure 11, the synthetic images were generated assuming a staircase like three-dimensional structure. The steps of the staircase have textures that are the same as those used in Figure 9. The computed depth map is again very accurate. The depth discontinuities are sensed with sharpness preserved, demonstrating the high spatial resolution of the proposed algorithm. Spikes in the two left strips are again due to extremely low depth confidence values in these areas. In the case of natural textures with enough texture contrast, such low confidence values are unlikely as other frequencies in the texture will provide sufficient information for robust depth estimation. (a) gray-coded depth map (a) wire frame plot of the depth map Figure 11: Depth from defocus applied to synthetic images of a staircase. The textures of the stairs are the same as those of the strips in Figure 9. The depth discontinuities are estimated with high accuracy reflecting high spatial resolution produced by the proposed algorithm. 5.2 Experiments on Real Images Images of real scenes were taken using a SONY XC-77 monochrome camera. The lens used is a Cosmicar B1214D-2 with f=25mm. The lens was converted into a telecentric lens by using an additional aperture to make its magnification invariant to defocus (see appendix and [Watanabe and Nayar-1995b]). As a result of telecentricity, image shifts between the far and near focused images are lower than 1/10 of a pixel. The lens aperture was set to F/8.3. The far-focused image i 1 was taken with the lens focused at 869mm from the camera, and the near-focused image i 2 with the lens focused at 529mm. These two distances were chosen so that all scene points lie between them. The above focus settings result in a maximum blur circle radius of e=F pixels. For each of the two focus settings, 256 images were averaged over 8.5 sec to get images with high signal-to-noise ratio. Figure 12 shows results obtained for a scene that includes a variety of textures. Figure 12(a) and (b) are the far-focused and near-focused images, respectively. Figure 12(c) and (d) are the computed depth map and its wireframe plot. Depth maps of all the curved and planar surfaces are detected with high fidelity and high resolution without any post-filtering. After 9\Theta9 median filtering, we get an even better depth map as shown in Figure 12(e). (a) far-focused image (b) near-focused image (c) gray-coded depth map without post-filtering (d) wireframe plot of (c) wireframe plot after 9\Theta9 median filtering Figure 12: The depth from defocus algorithm applied to a real scene with complex textures. Figure 13 shows results for a scene which includes areas with extremely weak textures, such as, the white background and the clay cup. Figure 13(a) and (b) are the far-focused and near-focused images, respectively. Figure 13(c) and (d) are the computed depth map and its wireframe plot. All image areas, except the white background area, produce accurate depth estimates. The depth confidence value oe in the textured background is 0.5% of the object distance. 11 The error on the table surface is 1.0% relative to object distance. Even the white background area has a reasonable depth map despite the fact that its texture is very weak. We see that the confidence map in Figure 13(e) reflects this lack of texture. This has motivated us to develop a modified algorithm, called adaptive coefficient smoothing, that repeatedly averages the coefficients computed by the rational operators until the confidence value reaches a certain acceptable level. Figure 13(f) shows the depth map computed using this algorithm. The last experiment seeks to quantify the accuracy of depth estimation. The target used is a plane paper similar to the textured background in the scene in Figure 12. This plane is moved in steps of 25mm and a depth map of the plane is computed for each position. Since the estimated depth fi is measured on the image side, it is mapped to the object side using the lens law of equation (2). The optical settings and processing conditions are the same as those used in the previous experiments. The plot in Figure 14 illustrates that the algorithm has excellent depth estimation linearity. The RMS error of a line fit to the measured depths is 4.2 mm. The slight curvature of the plot is probably due to errors in optical settings, such as, focal length and aperture. Depth values for a 50\Theta50 area were used to estimate the RMS depth error for each position of the planar surface. In Figure 14 the RMS errors are plotted as \Sigmaoe error bars. The RMS error relative to object distance is seen to vary with object distance. It is 0.4% - 0.8% for close objects and 0.8% - 1.2% for objects farther than 880 mm. This is partly because of the mapping from the depth measured on the image side to depth on the object side. The other reason is that the error in estimated depth oe fi is larger for a scene point with larger fi, as seen from equation (34). Note that this RMS error depends on the coefficient smoothing and post-filtering stages. We found empirically that the error has a Gaussian-like distribution. Using this distribution, one can show that the error reduces by a factor of 1/8 if the depth map is convolved with an 8\Theta8 averaging filter. This definition of error is often used to quantify the performance of range sensors. (a) far-focused image (b) near-focused image (c) gray-coded depth map (d) wireframe of the depth map (confidence value) 1=2 map (f) wireframe after adaptive coefficient smoothing Figure 13: Depth from defocus applied to a scene that includes very weak texture (white background). The larger errors in the region of weak texture is reflected by the confidence map. An adaptive coefficient smoothing algorithm uses the confidence map to refine depth estimates in regions with weak texture. actual depth (mm) estimated depth fitting error Figure 14: Depth estimation linearity for a textured plane. The plane is moved in increments of 25mm, away from the lens. All plotted distances are measured from the lens. The RMS error relative to object distance is 0.4% - 1.2%. 6 Conclusions We proposed the class of rational operators for passive depth from defocus. Though the operators are broadband, when used together, they provide invariance to scene texture. Since they are broadband, a small number of operators are sufficient to cover the entire frequency spectrum. Hence, rational operators can replace large filter banks that are expensive from a computational perspective. This advantage comes without the need to sacrifice depth estimation accuracy and resolution. We have detailed the procedure used to design rational operators. As an example, we constructed 7\Theta7 operators using a polynomial model for the normalized image ratio. However, the notion of rational operators is more general and represents a complete class of filters. The design procedure described here can be used to construct operators based on other rational models for the normalized image ratio. Further, rational operators can be derived for any desired blur function. In addition to the rational operators, we discussed a wide range of issues that are pertinent to depth from defocus. In particular, detailed analyses and techniques were provided for prefiltering near and far focused images as well as post-processing the outputs of the rational operators. The operator outputs have also been used to derive a depth confidence measure. This measure can be used to enhance computed depth maps. The proposed depth from defocus algorithm requires only a total of 5 convolutions. We tested the algorithm using both synthetic scenes and real scenes to evaluate performance. We found the depth detection gain error to be less than 1%, regardless of texture frequency. Depth accuracy was found to be 0:5 - 1:2% of object distance from the sensor. These results have several natural extensions. (a) Since some scene areas are expect to have very low texture frequency, it would be meaningful to embed the proposed scheme in a pyramid-based processing framework. Image areas with dominant low frequencies will have higher frequencies at higher levels of the pyramid. The proposed algorithm can be applied to all levels of the pyramid and the resulting depth maps can be combined using the depth confidence measures. (b) Given the efficiency of the al- gorithm, it is worth implementing a real-time version using a pipeline image processing architecture such as the Datacube MV200. We estimate that such an algorithm would result in at least 6 depth maps per second of 512\Theta480 resolution. (c) In our present im- plementation, we have varied the position of the image sensor to change the focus setting. Alternatively, the aperture size can be varied. Rational operators can be derived for such an optical setup using the basis functions b P1 (see [Watanabe and Nayar-1995a]). (d) Finally, it would be worthwhile applying the algorithm to outdoor scenes with large structures. Appendix 7.1 Problem of Image Registration For the rational operators to give accurate results, the far-focused image i 1 and near- focused image i 2 need to be precisely registered (within 0.1 pixel) with respect to one another. However, in most conventional lenses, magnification varies with focus setting and hence misregistration is introduced. Further, in our experiments, we have mechanically changed the focus setting and, in the process, introduced some translation between the two images. If the lens aberrations are small, the misregistration is decomposed into two factors - a global magnification change and a global translation. Of the two factors, magnification change proves much more harmful. This change can be corrected using image warping techniques [Darrell and Wohn-1988, Wolberg-1990]. However, this generally introduces undesirable effects such as smoothing and aliasing since warping is based on spatial interpolation and resampling techniques. We have used an optical solution to the problem that is described in the following section and detailed in [Watanabe and Nayar-1995b]. 7.1.1 Telecentric Optics In the imaging system shown in Figure 1, the effective image location of point P moves along the principal ray R as the sensor plane is displaced. This causes a shift in image coordinates of the image of P . This variation in image magnification with defocus manifests as a correspondence like problem in depth from defocus, as corresponding points in images are needed to estimate blurring. We approach the problem from an optical perspective rather than a computational one. Consider the image formation model shown in Figure 15. The only modification made with respect to the model in Figure 1 is the use of the external aperture A 0 . The aperture is placed at the front-focal plane, i.e. a focal length in front of the principal point O of the lens. This simple addition solves the problem of magnification variation with distance ff of the sensor plane from the lens. Simple geometrical analysis reveals that a ray of light R 0 from any scene point that passes through the center O 0 of aperture A 0 emerges parallel to the optical axis on the image side of the lens [Kingslake-1983]. As a result, despite blurring, the effective image coordinates of point P in both images i 1 are the same as the coordinate of its focused image Q on i f . Given an off-the- shelf lens, such an aperture is easily appended to the casing of the lens. The resulting optical system is called a telecentric lens. While the nominal and effective F-numbers of the classical optics in Figure 1 are f/a and d i /a, respectively, they are both equal to f/a 0 in the telecentric case. The magnification change can be reduced to an order of less than 0.03%, i.e. 0.1 pixel for a 640\Theta480 image. A detailed discussion on telecentricity and its implementation can be found in [Watanabe and Nayar-1995b]. We recently used this idea to develop a real-time active depth from defocus sensor [Nayar et al.-1995, Watanabe et al.-1995]. 7.1.2 Translation Correction We have seen in the previous section how magnification changes between the far-focused and near-focused images can be avoided. When the focus setting is changed, translations may also be introduced. Translation correction can be done using image processing without introducing any harmful image artifacts. However, the processing must be carefully R a' O ' f f a Figure 15: A constant-magnification imaging system for depth from defocus is achieved by simply placing an aperture at the front-focal plane of the optics. The resulting telecentric optics avoids the need for registering the far-focused and near-focused images [ Watanabe and Nayar-1995b implemented since we seek 0.1 pixel registration between the two images. The procedure we use is briefly described here and is detailed in [Watanabe and Nayar-1995b]. We use FFT-phase based local shift detection to estimate shift vectors with sub-pixel accuracy. We divide the Fourier spectra of corresponding local areas of the two images. Then we fit a plane to the phases of the ratio of the spectra. The gradient of the fitted plane is nothing but the relative shift between the two images. Once we get shift vectors at several positions in the image, similarity transform is used to model the shift vector field. By fitting the vectors to the similarity model, we can estimate the global translation and any residual magnification changes, separately [Watanabe and Nayar-1995b]. The residual magnification is corrected by tuning the aperture position of the telecentric optics. The translation is corrected by shifting both images in opposite directions. As we need sub-pixel accuracy, we interpolate the image and resample it to generate the registered images. The interpolating function is the Lanczos4 windowed sinc function [Wolberg- 1990]. Since the translation correction remains constant over the entire image, a single shift invariant convolution achieves the desired shift. Though this convolution distorts the image spectrum, since both images undergo the same amount of shift, the distortion is the same for both images. This common distortion is eliminated when the normalized image ratio M=P is computed before the application of the rational filters. After the above translation correction, we found the maximum registration error in our experiments to be as small as 0.02 pixels. 7.2 Operator Response and Depth Error The deviation of the ratio functions G P i (u; v) or GM i (u; v) after filter design, to those obtained from fitting the polynomial model to the normalized image ratio, varies with frequency (u; v), and hence depends on the texture of the scene. For the filter design described in section 4.1, we need a relation between the above ratio error and the depth estimation error. Starting with equation (12), we get: G P1 (u; v) G P2 (u; v) Here, fi f (u; v) is the depth estimated at a single frequency (u; v). Since -(u; v) in the ratio condition (13) has not been fixed, we can define G P1 (u; becomes: By differentiation we get: where, M(u; v; ff) and P (u; v; ff) can be treated as constants since we wish to find the error in fi f (u; v) caused by errors in GM1 (u; v) and G P2 (u; v). Solving for dfi f (u; v), we get: Since G P2 (u; v) is a small correction factor, it can be approximated by: From equations (23) and (20), since c P2 - c P1 (c P2 represents a small correction), the depth estimate fi can be approximated by integrating over all frequencies: du dv Hence, the error in fi caused by the error in fi f (u; v) is: Combining this expression with equation (43), we have: du dv What are the optimal values of dGM1 (u; v) and dG P2 (u; v) that would minimize the depth error dfi? This question is not trivial as dGM1 (u; v) and dG P2 (u; v) influence each other in a complex way. To avoid either of the two terms in the integrand in the numerator from taking on a disproportionately large value, we have decided to assume both terms to be constant of value -. This gives us the following bounds on dGM1 (u; v) and dG P2 (u; v): oe GM1 (u; oe G P2 (u; where, the jfi f (u; v)j was set to 1 as this represents the worst case, i.e. largest normalized depth error. Acknowledgements This research was conducted at the Center for Research in Intelligent Systems, Department of Computer Science, Columbia University. It was supported in part by the Production Engineering Research Laboratory, Hitachi, and in part by the David and Lucile Packard Fellowship. The authors thank Yasuo Nakagawa of Hitachi Ltd. for his support and encouragement of this work. --R Range imaging sensors. Principles of Optics. The Fourier Transform and Its Applications. The Laplacian pyramid as a compact image code. Pyramid based depth from focus. A matrix based method for determining depth from focus. Computing depth from out-of-focus blur using a local frequency representation Introduction to Mathematical Statistics. Focusing. Memo 160 Robot Vision. A perspective on range finding techniques for computer vision. A Pyramid Framework for Early Vision. Optical System Design. Journal of Computer Vision Shape from focus: An effective approach for rough surfaces. The Science of Fractal Images. A new sense for depth of field. Numerical Recipes in C. Depth from defocus: A spatial domain approach. Parallel depth recovery by changing camera parameters. Minimal operator set for texture invariant depth from defocus. Telecentric optics for constant-magnification imaging Digital Image Warping. Depth from focusing and defocusing. Moment filters for high precision computation of focus and stereo. --TR --CTR M. Boissenin , J. Wedekind , A. N. Selvan , B. P. Amavasai , F. Caparrelli , J. R. Travis, Computer vision methods for optical microscopes, Image and Vision Computing, v.25 n.7, p.1107-1116, July, 2007 A. N. Rajagopalan , S. Chaudhuri , Uma Mudenagudi, Depth Estimation and Image Restoration Using Defocused Stereo Pairs, IEEE Transactions on Pattern Analysis and Machine Intelligence, v.26 n.11, p.1521-1525, November 2004 Liming Chen , Georgy Kukharev , Tomasz Ponikowski, The PCA reconstruction based approach for extending facial image databases for face recognition systems, Enhanced methods in computer security, biometric and artificial intelligence systems, Springer-Verlag, London, 2005 M. Boissenin , J. Wedekind , A. N. Selvan , B. P. Amavasai , F. Caparrelli , J. R. Travis, Computer vision methods for optical microscopes, Image and Vision Computing, v.25 n.7, p.1107-1116, July, 2007 Vinay P. Namboodiri , Subhasis Chaudhuri, On defocus, diffusion and depth estimation, Pattern Recognition Letters, v.28 n.3, p.311-319, February, 2007 Paolo Favaro , Stefano Soatto, A Geometric Approach to Shape from Defocus, IEEE Transactions on Pattern Analysis and Machine Intelligence, v.27 n.3, p.406-417, March 2005 Reinhard , Erum Arif Khan, Depth-of-field-based alpha-matte extraction, Proceedings of the 2nd symposium on Applied perception in graphics and visualization, August 26-28, 2005, A Coroa, Spain Lyubomir Zagorchev , Ardeshir Goshtasby, A paintbrush laser range scanner, Computer Vision and Image Understanding, v.101 n.2, p.65-86, February 2006

scene textures;texture invariance;depth estimation;blur function;depth confidence measure;normalized image ratio;real-time performance;passive depth from defocus;broadband rational operators

294178

The Efficient Computation of Sparse Jacobian Matrices Using Automatic Differentiation.

This paper is concerned with the efficient computation of sparse Jacobian matrices of nonlinear vector maps using automatic differentiation (AD). Specifically, we propose the use of a graph coloring technique, bicoloring, to exploit the sparsity of the Jacobian matrix J and thereby allow for the efficient determination of J using AD software. We analyze both a direct scheme and a substitution process. We discuss the results of numerical experiments indicating significant practical potential of this approach.

Introduction The efficient numerical solution of nonlinear systems of algebraic equations, F usually requires the repeated calculation or estimation of the matrix of first derivatives, the Jacobian In large-scale problems matrix J is often sparse and it is important to exploit this fact in order to efficiently determine, or estimate, matrix J at a given argument x. This paper is concerned with the efficient calculation of sparse Jacobian matrices by the judicious application of automatic differentiation techniques. Specifically, we show how to define "thin" matrices V and W such that the nonzero elements of J can easily be extracted from the calculated pair (W T J; JV ). This research was partially supported by the Applied Mathematical Sciences Research Program (KC-04-02) of the Office of Energy Research of the U.S. Department of Energy under grant DE-FG02-90ER25013, and in part by the Advanced Computing Research Institute, a unit of the Cornell Theory Center which receives major funding from the National Science Foundation and IBM Corporation, with additional support from New York State and members of its Corporate Research Institute. This technical report also appears as Cornell Computer Science Technical Report TR 95-1557. y Computer Science Department and Center for Applied Mathematics, Cornell University, Ithaca NY 14850. z Computer Science Department, Cornell University, Ithaca NY 14850. Given an arbitrary n-by-t V matrix V , product JV can be directly calculated using automatic differentiation in the "forward mode"; given an arbitrary m-by-t W matrix W , the product W T J can be calculated using automatic differentiation in the "reverse mode", e.g., [11, 13]. The forward mode of automatic differentiation allows for the computation of product JV in time proportional to t V \Delta !(F ) where !(F ) is the time required to evaluate F . This fact leads to the following practical question. Given the structure of a sparse Jacobian matrix J , how can a matrix V be chosen so that the nonzeros of J can easily be determined from the product JV ? A good solution is offerred by the sparse finite-differencing literature [4, 5, 6, 7, 8, 10] and adapted to the automatic differentiation setting [1]. Partition the columns of J into a set of groups GC , where the number of groups in GC is denoted by jGC j, such that the columns in each group G 2 GC are structurally orthogonal 1 . Each group G 2 GC defines a column v of only if column i is in group G; otherwise, v It is clear that the nonzeros of J can be immediately "identified" from the computed product JV . Graph coloring techniques, applied to the column intersection graph of J , can be used to try and produce a partition GC with low cardinality jGC j. This, in turn, induces a thin matrix V , i.e., construct However, it is not always possible to ensure that jGC j is small: consider a sparse matrix J with a single dense row. Alternatively, the reverse mode of automatic differentiation allows for the computation of the product W T J in time proportional to t W \Delta !(F ) where !(F ) is the time required to evaluate F , and . The "transpose" of the argument above can lead to an efficient way to determine J . That is, apply graph coloring techniques to the row intersection graph of J to induce a thin matrix via the reverse mode of automatic differentiation (takes time proportional to trivially extract the nonzeros of J from the computed matrix W T J . Of course it is easy to construct examples where defining a thin matrix W is not possible - consider the case where J has a dense column. Clearly there are problems where a row-oriented approach is preferable, there are problems where a column-oriented approach is better. Unfortunately, it is easy to devise problems where neither approach is satisfactory: let J have both a dense row and a dense column. This is exactly when it may pay to use both modes of automatic differentiation simultaneously: compute a pair choices of W and V , and extract the nonzero elements of J from this computed pair of thin matrices. Our concern in this paper is with efficiency with respect to number of floating point operations or f lops. We do not concern ourselves with space requirements in this study. However, it should be noted that the reverse mode of automatic differentiation often requires significantly more space than the forward mode: if space is tight then our suggested approach, which involves application of both forward and reverse modes, may not be possible. There is current research activity on reducing the space requirements of the reverse mode of automatic differentiation, e.g., [12]. We note that an independent proposal regarding sparse Jacobian calculation is made by Hossain and Steihaug [15]: a graph-theoretic interpretation of the direct determination problem is given and an algorithm based on this interpretation is provided. In this paper we proffer a new direct method and we also propose a substitution method, both based directly on the Jacobian structure. We compare our direct and substitution methods, numerically, and we discuss the round-off properties of the substitution method. In addition, we interface our graph coloring software to the automatic differentiator ADOL-C, [14], and report on a few preliminary computational results. The remainder of the paper is organized as follows. In x2, we review the relevant aspects of Two nonzero n-vectors v; w are structurally orthogonal if v i w automatic differentiation, both forward and reverse modes. In x3, we formalize the combinatorial problems to be solved both from a matrix point of view and in terms of graph theory. We propose both a direct determination problem and a substitution problem. In x4, we propose "bi-coloring" approaches to both the direct determination and "determination by substitution" problems. The bi-coloring technique produces matrices V and W where JV is subsequently determined via the forward mode of automatic differentiation, W T J is detemined via the reverse mode. Typically, the column dimensions of V and W will be small: the cost of the application of automatic differentiation is proportional to the sum of the column dimensions of V and W (times the work to evaluate F ). In x5 we present and discuss various numerical experiments. The experiments indicate that our bi-coloring approach can significantly reduce the cost of determining J (over one-sided Jacobian determination). The substitution method we propose consistently outperforms the direct method. However, the substitution calculation increases the chance of round-off contamination. This effect is discussed in x6. We end the paper in x7 with some concluding remarks and observations on possible directions for future research. Specifically, we note that while sparsity is a symptom of underlying structure in a nonlinear problem, it is not a necessary symptom. Moreover, it is often possible to exploit structure in the absence of sparsity and apply AD tools "surgically" to efficiently obtain the Jacobian matrix J . /sectionBasics of automatic differentiation Automatic differentiation is a chain rule based technique for evaluating the derivatives analytically (and hence without any truncation errors) with respect to input variables of functions defined by a high-level language computer program. In this section we briefly review the basics of automatic differentiation, borrowing heavily from [11, 13]. A program computing the function can be viewed as a sequence of scalar assignments v where the vector v can be thought of as set of ordered variables such that v i computed before v j using the set of variables fv k jk ! ig. Here / j represent elementary functions, which can be arithmetic operations and/or univariate transcedental functions. Ordering the variables as above, we can partition variables v j into three vectors. In general, the number of intermediate variables is much larger than the dimensions of the problem, Assume that all these elementary functions / i are well defined and have continuous elementary Assuming with out loss of generality that the dependent variables y do not themselves occur as arguments of elementary functions, we can combine the partials c ij into the (p +m) \Theta (n + p) matrix c n+i;j 1-i-p+m Unless elementary functions with more than two arguments are included in the library, each row of C contains either one or two non zero entries. We define a number Also, since work involved in an elementary function is proportional to number of arguments, it follows that proportionality. Because of the ordering relation the square matrix C is upper trapezoidal with nontrivial superdiagonals. Thus C can be partitioned as A L where and L is strictly lower triangular. Application of the chain rule yields: !/ \Deltax \Deltay If we eliminate the intermediate vector \Deltay from above (1.4), we get an expression(the Schur complement) for the Jacobian: is a unit lower triangular matrix, the calculation of the matrix products ~ leads to two natural ways to compute J : A or The alternative expressions for J given in (1.6) define the two basic modes of automatic differenti- ation, forward and reverse. The forward mode corresponds to computing the rows of ~ A, one by one, as the corresponding rows of [A; are obtained from successive evaluation of elementary functions. Since this amounts to solutions of n linear systems with lower-triangular matrix [I \Gamma L], followed by multiplication of dense columns of ~ A by M , the total computational effort is roughly n \Delta q or n \Delta !(F ). The reverse mode corresponds to computing ~ M as solution to linear system M T . This back-substitution process can begin only after all elementary functions and their partial derivatives have been evaluated. Since this amounts to the solution of m linear systems with lower triangular matrix [I \Gamma L], followed by multiplication of dense rows of ~ M by A, total computational effort is roughly m \Delta q or m \Delta !(F ). We are interested in computing products of the form JV and W T J . Product JV can be computed: which can clearly be done in time proportional to t V n\Thetat V . Analagously, product can be computed: which can be done in time proportional to t W \Delta !(F ) assuming An important subcase worthy of special attention is when F is a scalar function, i.e., In this case the Jacobian matrix corresponds to the transpose of the gradient of F and is a single row vector. Note that the complexity arguments applied to this case imply that the reverse mode of AD yields the gradient in time proportional to !(F ) whereas the forward mode costs n \Delta !(F ). The efficiency of the forward mode evaluation of the gradient can be dramatically increased - i.e., the dependence on n is removed - if F has structure that can be exploited [2]. Partition problems and graph theory Our basic task is to efficiently determine thin matrices V; W so that the nonzero elements of J can be readily extracted from the information (W T J; JV ). The pair of matrices (W T J; JV ) is obtained from the application of both modes of automatic differentiation: matrix W T J is computed by the reverse mode, the forward mode determines JV . The purpose of this section is to more rigorously formulate the question of determining suitable matrices V; W , first in the language of "partitions", and then using graph theoretic concepts. We begin with an example illustrating the usefulness of simultaneously applying both modes of AD, forward and reverse. Consider the following n-by-n Jacobian, symmetric in structure but not in value: It is clear that a partition of columns consistent with the direct determination of J requires n groups. This is because a "consistent column partition" requires that each group contain columns that are structurally orthogonal and the presence of a dense row implies each group consists of exactly one column. Therefore, if matrix V corresponds to a "consistent column partition" then V has n columns and the work to evaluate JV by the forward mode of AD is proportional to n \Delta !(F ). By a similar argument, and the fact that a column of J is dense, a "consistent row partition" requires n groups. Therefore, if matrix W corresponds to a "consistent row partition" then W has rows and the work to evaluate W T J by the reverse mode of AD is proportional to n \Delta !(F ). Definition 2.1 A bi-partition of a matrix A is a pair (GR ; GC ) where GR is a row partition of a subset of the rows of A , GC is a column partition of a subset of the columns of A. In this example the use of a bi-partition dramatically decreases the amount of work required to determine J . Specifically, the total amount of work required is proportional to 3 \Delta !(F ). To see this the usual convention of representing the i th column of the identity matrix with e i . Clearly elements are directly determined from the product JV ; elements 4 are directly determined from the product W T J . The basic idea is to partition the rows into a set of groups GR and the columns into a set of groups GC , with small as possible, such that every nonzero element of J can be directly determined from either a group in GR or a group in GC . Definition 2.2 A bi-partition (GR ; GC ) of a matrix A is consistent with direct determination if for every nonzero a ij of A, either column j is in a group of GC which has no other column having a nonzero in row i, or row i is in a group of GR which has no other rows having a nonzero in column j. Clearly, given a bi-partition (GR ; GC ) consistent with direct determination, we can trivially construct matrices n\ThetajG C j such that A can be directly determined from If we relax the restriction that each nonzero element of J be determined directly then it is possible that the work required to evaluate the nonzeroes of J can be further reduced. For example we could allow for a "substitution" process when recovering the nonzeroes of J from the pair Figures 2.1, 2.2 illustrate that a substitution method can win over direct determination: Figure 2.1 corresponds to direct determination, Figure 2.2 corresponds to determination using Figure 2.1: Optimal partition for direct method In both cases elements labelled are computed from the column grouping, i.e., calculated using the product JV ; elements labelled 4 are calculated form the row groupings, i.e., calculated using the product W T J . The matrix in Figure 2.1 indicates that we can choose GC with jGC GR with jGR determine all elements directly. That is, choose choose Therefore in this case the work to compute J satisfies !(J) - 3 \Delta !(F ). Note that some elements can be determined twice, e.g., J 11 . However, the matrix in Figure 2.2 shows how to obtain the nonzeroes of J , using substitution, in work proportional to 2 \Delta !(F ). Let be the (forward) computed Figure 2.2: Optimal partition for substitution method vector 2 be the (reverse) computed row vector p T Most of the nonzero elements are determined directly (no conflict). The remaining elements can be resolved, It is easy to extend this example so that the difference between the number of groups needed, between substitution and direct determination, increases with the dimension of the matrix. For example, a block generalization is illustrated in Figure 2.3: if we assume l ? 2w it is straightforward to verify that in the optimal partition the number of groups needed for direct determination will be 3w and determination by substitution requires 2w groups. Definition 2.3 A bi-partition (GR ; GC ) of a matrix A is consistent with determination by substitution, if there exists an ordering - on elements a ij , such that for every nonzero a ij of A, either column j is in a group where all nonzeros in row i, from other columns in the group, are ordered lower than a ij , or row i is in a group where all the nonzeros in column j, from other rows in the group, are ordered lower than a ij . In the usual way we can construct a matrix V from the column grouping GC and a matrix W from the row grouping GR : for example, to construct the columns of V associate with each group l l Figure 2.3: Block example in GC a boolean vector, with unit entries indicating membership of the corresponding columns. We can now state our main problem(s) more precisely: The bi-partition problem (direct) : Given a matrix A, obtain a bi-partition (GR ; GC ) consistent with direct determination, such that total number of groups, minimized. The bi-partition problem (substitution) : Given a matrix A, obtain a bi-partition (GR ; GC ) consistent with determination by substitution, such that total number of groups, jGC j, is minimized. The bi-partition problems can also be expressed in terms of graphs and graph coloring. This graph view is important in that it more readily exposes the relationship of the bi-partition problems with the combinatorial approaches used in the sparse finite-differencing literature, e.g., [4, 5, 6, 7, 8]. However, we note that the remainder of this paper, with the exception of the error analysis in x6, does not rely directly on this graph interpretation. To begin, we need the usual notion of a coloring of the vertices of a graph, the definition of a bipartite graph, and the concept of path coloring [4, 7] specialized to the bipartite graph case. A p-coloring of a graph is the set of vertices or nodes and E is the set of edges, is a function such that OE(u) 6= OE(v), if (u; v) 2 E. The chromatic number -(G) is the smallest p for which G has a p-coloring. A p-coloring OE of G induces a partition of vertices such that Given a matrix A 2 ! m\Thetan , define a bipartite graph E) where corresponds to the jth column of A, and r i corresponds to the ith row of A. There is an edge, (r a nonzero in A. In [4, 7] a path p-coloring of a graph is defined to be a vertex coloring using p colors with the additional property that every path of at least 3 edges uses at least 3 colors. Here we need a slight modification of that concept appropriate for the direct determination problem. We note that "color 0" is distinguished in that it corresponds to the lack of a true color assignment: i.e., indicates that vertex i is not assigned a color. Definition 2.4 Let G E) be a bipartite graph. A mapping OE is a bipartite path p-coloring of G b if 1. Adjacent vertices have different assignments, i.e., if (i; 2. The set of positive colors used by vertices in V 1 is disjoint from the set of positive colors used by vertices in 0g. 3. If vertices i and j are adjacent to vertex k with 4. Every path of 3 edges uses at least at least 3 colors. The smallest number for which graph G b is bipartite path p-colorable is denoted by - Figure 2.4 shows a valid bipartite path p-coloring. Numbers adjacent to the vertices denote colors. We note that Hossain and Steihaug [15] define a similar concept. However, their definition of path p-coloring does not allow for the "uncolor assignment", i.e., Consequently, a technique to remove empty groups is needed [15]. We are now in position to state the graph analogy to the concept of a bi-partition consistent with direct determination. Theorem 2.1 Let A be a m \Theta n matrix with corresponding bipartite graph E). The mapping induces a bi-partition (GR ; GC ), with consistent with direct determination if and only if OE is a bipartite path p-coloring of G b (A). Proof. (() Assume that OE is a bipartite path p-coloring of G b (A), inducing a bi-partition (GR ; GC ) of rows and columns of A. If this bi-partition is not consistent with direct determination, then there is a nonzero element a ij in the matrix for which the definition "either column j is in a group of GC which has no other column having a nonzero in row i, or row i is in a group of GR which has no other rows having a nonzero in column j" doesn't hold. This can happen only if one of the following cases hold: there exists a column q with a iq 6= 0, such that OE(c j this contradicts condition 3. there exists a row p with a pj 6= 0, such that OE(r i contradicts condition 3. COLUMNS ROWS000000 Figure 2.4: A valid bipartite path coloring There exists a column q and a row p, such that columns j and q are in the same group with a iq 6= 0 and rows i and p are in the same group with a pj 6= 0 This implies OE is a 2-coloring of the path (r contradiction of condition 4. Conversely, assume that OE induces a bi-partition consistent with direct determination of A. It is clear that conditions 1 \Gamma 3 must be satisfied. It remains for us to establish condition 4: i.e., every path of 3 edges uses at least 3 colors. Suppose there is a bi-colored by condition 3 the two colors on this path are positive. It is easy to see that element a jk cannot be determined directly: there is a conflict in row group there is a conflict in column group OE(c k these are the only two chances to determine a jk . . To capture the substitution notion the cyclic p-coloring definition [4] is modified slightly and applied to a bipartite graph. Definition 2.5 Let A be a m \Theta n matrix with corresponding bipartite graph E). A mapping OE is a bipartite cyclic p-coloring of G b if 1. Adjacent vertices have different assignments, i.e., if (i; 2. The set of positive colors used by vertices in V 1 is disjoint from the set of positive colors used by vertices in 0g. 3. If vertices i and j are adjacent to vertex k with 4. Every cycle uses at least at least 3 colors. The smallest number for which graph G b is bipartite cyclic p-colorable is denoted by - c (G b ). Figure 2.5 shows a valid bipartite cyclic p-coloring note that only 2 colors are necessary whereas the bipartite path p-coloring in Figure 2.4 requires 3 colors. The notion of a bi-partition consistent12COLUMNS Figure 2.5: A valid bipartite cyclic coloring with determination via substitution can now be cleanly started in graph-theoretic terms. Theorem 2.2 Let A be a m \Theta n matrix with corresponding bipartite graph E). The mapping OE : induces a bi-partition (GR ; GC ), with consistent with determination by substitution if and only if OE is a bipartite cyclic p-coloring of G b (A). Proof. ()) Assume OE induces a bi-partition consistent with determination by substitution but OE is not a bipartite cyclic p-coloring of G b (A). Clearly condition must hold; it is easy to see that if condition 3 doesn't hold then not all nonzero elements can be determined. The only non-trivial violation is condition 4: there is a cycle which has only two colors, i.e all the vertices in the cycle have the same color c 1 , and all the vertices in the cycle have the same color c 2 . Note that neither c 1 and c 2 can be equal to 0, since a node colored 0 in a cycle would imply that its adjacent vertices are both colored differently, implying that there are at least 3 colors. Consider the submatrix A s of A, corresponding to this cycle. Submatrix A s has at least two non zeros in each row and in each column, since each vertex has degree 2 in the cycle. But since we are considering substitution methods only, at least one element of A s needs to be computed directly. Clearly there is no way to get any element of this submatrix directly, a contradiction. Conversely, assume that OE is a bipartite cyclic p-coloring of G b (A) but that bi-partition induced by OE is not consistent with determination by substitution. But, edges (nonzeros) with one end assigned color "0" can be determined directly: by the definition of bi-coloring there will be no conflict. Moreover, every pair of positive colors induces a forest (i.e., a collection of trees); therefore, the edges (nonzeros) in the induced forest can be resolved via substitution [4]. The two bi-partition problems can now be simply stated in terms of optimal bipartite path and cyclic p-colorings: The bipartite path p-coloring problem : Determine a bipartite path p-coloring of G b (A) with the smallest possible value of p, i.e., The bipartite cyclic p-coloring problem : Determine a bipartite cyclic p-coloring of G b (A) with the smallest possible value of p, i.e., The graph theoretic view is useful for both analyzing the complexity of the combinatorial problem and suggesting possible algorithms, exact or heuristic. In fact, using the the p-coloring notions discussed above, and an approach similar to that taken in [4], it is easy to show that corresponding decision problems are NP-complete. Bipartite cyclic p-coloring decision problem (CCDP): Given an integer p - 3 and an arbitrary bipartite graph G, is it possible to assign a cyclic p-coloring to nodes of G? Bipartite path p-coloring decision problem (PCDP) : Given an integer p - 3 and an arbitrary bipartite graph G, is it possible to assign a bipartite path p-coloring to nodes of G? The proofs are a straightforward adaptation of those in [4] and we omit them here. The upshot of these (negative) complexity results is that in practise we must turn our attention to (fast) heuristics to approximately solve the cyclic and path coloring problems. In the next section we present simple, effective, and "easy-to-visualize" heuristics for these two combinatorial problems. Finally, it is easy to establish a partial ordering of chromatic numbers: where G(M) refers to the column intersection graph of matrix M , -(G(M)) is the (usual) chromatic number of graph G(M ). The first inequality in (2.2) holds because if OE is a bipartite path p-coloring then OE is a bipartite cyclic p-coloring; the second inequality holds because a trivial way to satisfy conditions 1 \Gamma 4 of Definition 2.4 is to assign "0" to all the row (column) nodes and then use positive colors on all the column (row) nodes to satisfy condition 3. This ordering supports the tenet that use of bi- partition/bi-coloring is never worse than one-sided calculation and that a substitution approach is never worse than a direct approach (in principle). Bi-coloring The two combinatorial problems we face, corresponding to direct determination and determination by substitution, can both be approached in the following way. First, permute and partition the structure of J : ~ indicated in Figure 3.1. The construction of this partition is crucial; however, we postpone that discussion until after we illustrate its' utility. Assume I and R R Figure 3.1: Possible partitions of the matrix ~ Second, define appropriate intersection graphs G I R based on the partition [J C jJ R ]; a coloring of G I C yields a partition of a subset of the columns, GC , which defines matrix V . Matrix W is defined by a partition of a subset of rows, GR , which is given by a coloring of G I R . We call this double coloring approach bi-coloring. The difference between the direct and substitution cases is in how the intersection graphs, G I R , are defined, and how the nonzeroes of J are extracted from the pair (W T J; JV ). 3.1 Direct determination In the direct case the intersection graph G I C is defined: G I 9k such that J kr 6= 0; J ks 6= 0 and either (k; r) 2 JC or The key point in the construction of graph G I C , and why G I C is distinguished from the usual column intersection graph, is that columns r and s are said to intersect if and only if their nonzero locations overlap, in part, in JC : i.e., columns r and s intersect if J kr \Delta J ks = 0 and either The "transpose" of the procedure above is used to define G I R ). Specifically, G I I R if nnz(row If M is a matrix or a vector then "nnz(M)" is the number of nonzeroes in M In this case the reason graph G I R is distinguished from the usual row intersection graph is that rows r and s are said to intersect if and only if their nonzero locations overlap, in part, in JR : rows r and s intersect if J rk \Delta J The bi-partition (GR ; GC ), induced by coloring of graphs G I R and G I C , is consistent with direct determination of J . To see this consider a nonzero element (i; is in a group of GC (corresponding to a color) with the property that no other column in GC has a nonzero in row i: hence, element (i; j) can be directly determined. Analagously, consider a nonzero element row r will be in a group of GR (corresponding to a color) with the property that no other row in GR has a nonzero in column s: hence, element (r; s) can be directly determined. Since every nonzero of J is covered, the result follows. Example: Consider the example Jacobian matrix structure shown in Figure 3.2 with the partition 42 44 Figure 3.2: Example Partition The graphs GC and GR formed by the algorithm outlined above are given in Figure 3.3. Coloring GC requires 3 colors, while GR can be colored in two. Boolean matrices V and W can be formed in the usual way: each column corresponds to a group (or color) and unit entries indicate column (or row) membership in that group: \Theta J \Theta J 42 J 23 J 44 0 Clearly, all nonzero entries of J can be identified in either JV or W T J . c c 21212 r r r Figure 3.3: Graphs GC and GR (direct approach) 3.2 Determination by substitution The basic advantage of determination by substitution in conjunction with partition is that sparser intersection graphs G I R can be used. Sparser intersection graphs mean thinner matrices V ,W which, in turn, result in reduced cost. In the substitution case the intersection graph G I C is defined: G I 9k such that J kr 6= 0; J ks 6= 0 and both (k; r) 2 JC , (k; s) 2 JC . Note that the intersection graph G I captures the notion of two columns intersecting if there is overlap in nonzero structure in JC : columns r and s intersect if J kr \Delta J ks = 0 and both some k. It is easy to see that E I C is a subset of the set of edges used in the direct determination case. The "transpose" of the procedure above is used to define G I R ). Specifically, G I I R if row i 2 JR and nnz(row The intersection graph G I R ) captures the notion of two rows intersecting if there is overlap in nonzero structure in JR : rows r and s intersect if J rk \Delta J for some k. It is easy to see that E I R is a subset of the set of edges used in direct determination. All the elements of J can be determined from (W T J; JV ) by a substitution process. This is evident from the illustrations in Figure 3.4. Figure 3.4 illustrates two of four possible nontrival types of partitions. In both cases it is clear that nonzero elements in the section labelled "1" can be solved for directly - by the construction process they will be in different groups. Nonzero elements in "2" can either be determined directly, or will depend on elements in section "1". But elements in section "1" are already determined (directly) and so, by substitution, elements in "2" can be determined after "1". Elements in section "3" can then be determined, depending only on elements in "1" and "2", and so on until the entire matrix is resolved. JR Figure 3.4: Substitution Orderings Example. Consider again the example Jacobian matrix structure shown in Figure 3.2. Column and row intersection graphs corresponding to substitution are given in Figure 3.5. Note that GC is disconnected and requires 2 colors; GR is a simple chain and also requires 2 colors. c c 21112 r r r Figure 3.5: Graphs GC and GR for substitution process The coloring of GC and GR leads to the following matrices V , W and the resulting computation of JV , W \Theta 0 \Theta J \Theta J 42 J 23 J 44 0 It is now easy to verify that all nonzeroes of J can be determined via substitution. 3.3 How to partition J. We now consider the problem of obtaining a useful partition [J C jJ R ], and corresponding permutation matrices P ,Q, as illustrated in Figure 3.1. A simple heuristic is proposed based on the knowledge that the subsequent step, in both the direct and the substitution method, is to color intersection graphs based on this partition. Algorithm MNCO builds partition JC from bottom up, and partition JR from right to left. At the k th major iteration either a new row is added to JC or a new column is added to JR : the choice depends on considering a lower bound effect: ae(J T r where ae(A) is the maximum number of nonzeroes in any row of matrix A, r is a row under consideration to be added to JC , c is a column under consideration to be added to JR . Hence, the number of colors needed to color G I C is bounded below by ae(J C ); the number of colors needed to color G I R is bounded below by ae(J T In algorithm MNCO, matrix C) is the submatrix of J defined by row indices R and column indices C: M consists of rows and columns of J not yet assigned to either JC or JR . Minimum Nonzero Count Ordering (MNCO) 1. Initialize 2. Find r 2 R with fewest nonzeros in M 3. Find c 2 C with fewest nonzeros in M 4. Repeat Until if ae(J T R=R-frg else C=C-fcg repeat Note that, upon completion, JR ; JC have been defined; the requisite permutation matrices are implicitly defined by the ordering chosen in MNCO. Bi-coloring performance In this section we present results of numerical experiments. The work required to compute the sparse Jacobian matrix is the work needed to compute (W T J; JV ) which, in turn, is proportional to the work to evaluate the function F times the sum of the column dimensions of the boolean matrices . The column dimension sum, is equal to the number of colors used in the bi-coloring. In our experiments we compare the computed coloring numbers required for the direct and substitution approaches. We also compute the number of colors required by one-sided schemes: a column partition alone corresponds to the construction of V based on coloring the column intersection graph of J , a row partition alone corresponds to the construction of W based on coloring the row intersection graph of J . The latter case leads to the application of the reverse mode of AD (alone), whereas the former case leads to use of the forward mode. Both the direct and substitution methods require colorings of their respective pairs of intersection graphs, G I R . Many efficient graph coloring heuristics are available: in our experiments we use the incidence degree (ID) ordering [3, 8]. We use three sources of test matrices: a linear programming testbed with results reported in Table 1 and summarized in Table 2; the Harwell-Boeing sparse matrix collection, with results reported in Tables 3,4; self-generated m-by-n "grid matrices" with results given in Tables 5, 6. A grid matrix is constructed in the following way. First, approximately n of the columns are chosen, spaced uniformly. Each chosen column is randomly assigned DENS \Delta m nonzeroes. Second, approximately p m of the rows are chosen, spaced uniformly. Each chosen row is randomly assigned nonzeroes. We vary DENS as recorded in Table 5. For each problem we cite the dimensions of the matrix A and the number of nonzeros(nnz). The experimental results we report are the number of colors required by our bi-coloring approach, both direct and substitution, and the number of colors required by one-sided schemes. Bi-coloring One-sided Name m n nnz Direct Substitution column row standata stair 356 620 4021 36 29 36 36 blend 74 114 522 vtp.base 198 347 agg 488 615 2862 19 13 43 19 agg2 516 758 4740 26 21 49 43 agg3 516 758 4756 27 21 52 43 bore3d 233 334 1448 28 24 73 28 israel 174 316 2443 61 adlittle 56 138 Table 1: LP Constraint Matrices (http://www.netlib.org/lp/data/) Table 2: Totals for LP Collection Bi-coloring 1-sided Coloring Name M N NNZ Direct Substitution column row cannes 256 256 256 2916 cannes 268 268 268 1675 cannes cannes 634 634 634 7228 28 21 28 28 cannes 715 715 715 6665 22 cannes 1054 1054 1054 12196 31 23 cannes 1072 1072 1072 12444 chemimp/impcolc 137 137 411 6 4 8 9 chemimp/impcold 425 425 1339 6 5 11 11 chemimp/impcole 225 225 1308 21 14 chemwest/west0067 67 67 294 9 7 9 12 chemwest/west0381 381 381 2157 12 9 29 50 chemwest/west0497 497 497 1727 22 19 28 55 Table 3: The Harwell-Boeing collection (ftp from orion.cerfacs.fr) 4.1 Observations First, we observe that the bi-coloring approach is often a significant win over one-sided determi- nation. Occasionally, the improvement is spectacular, e.g., "cannes 715". Improvement in the Harwell-Boeing problems are generally more significant than on the LP collection in the sense that bi-coloring significantly outperforms both one-sided possibilities. This is partially due to the fact that the matrices in the LP collection are rectangular whereas the matrices in the Harwell-Boeing collection are square: calculation of the nonzeroes of J from W T J alone can be quite attractive when J has relatively few rows. The grid collection displays the advantage of bi-coloring to great effect - grid matrices are ideal bi-coloring candidates. In general the advantage of substitution over direct determination is not as great as the difference between bi-coloring and one-sided determination. Nevertheless, fewer colors are almost always needed and for expensive functions F this can be important. For most problems the gain is about 20% though it can approach 50%, e.g., "watt2". Table 4: Totals for Harwell-Boeing Collection Bi-coloring 1-sided coloring M N DENS Direct Substitution column row 100 100 0.52 20 20 84 74 100 100 0.64 20 20 95 93 100 100 1.00 20 20 100 100 100 400 0.53 100 400 0.64 100 400 1.00 Table 5: Grid Matrices 4.2 Interface with ADOL-C We have interfaced our coloring and substitution routines with the ADOL-C software. The C++ package ADOL-C [14] facilitates the evaluation of first and higher order derivatives of vector func- tion, defined by programs written in C or C++. We compare the time needed on a sample problem with respect to five approaches: ffl AD/bi-coloring (direct) ffl AD/bi-coloring (substitution) ffl AD/column coloring (forward mode) ffl AD/row coloring (reverse mode) ffl FD (Sparse finite differencing based on column coloring) The test function F we use is a simple nonlinear function: define be the index set of nonzeroes in row i of the Jacobian matrix and define F Table Totals for Grid Matrices The Jacobian matrix (and thus the sparsity pattern) is a 10-by-3 block version of Figure 2.3, i.e., 3. Problem dimensions as indicated in Figure 4.1, were used in the experiments. Our results, portrayed in Figure 4.1, suggest the following order of execution time requirement by different techniques: Note that FD requires more time than AD=column even though the same coloring is used for both. This is because the work estimate t V \Delta !(F ) is actually an upper bound on the work required by the forward mode where t V is the number of columns of V . This bound is often loose in practise finite-differencing since the subroutine to evaluate F is actually called times. Problem Size Time (in seconds) Performance graph of different sparse approaches AD Column AD Row FD AD-BiColoring-Direct AD-BiColoring-Substitution Figure 4.1: A comparison of different sparse techniques Another interesting observation is that the reverse mode calculation (AD/Row) is about twice as expensive as the forward calculation (AD/Column). This is noteworthy because in this example, based on the structure Figure 2.1, the column dimensions of V and W are equal. This suggests that it may be practical to weigh the cost of the forward calclulation of Jv versus the calculation of w T J , where w; v are vectors. We comment further on this aspect in x7. 5 Substitution and round-off In general, the substitution approach requires fewer colors and therefore is more efficient 3 , in principle, than direct determination. However, there is a possibility of increased round-off error due to the substitution process. In fact an analogous issue arises in the sparse Hessian approximation context [4, 7, 16] where, indeeed, there is considerable cause for concern. The purpose of this section is to examine this question in the AD context. The bottom line here is that there is less to worry about in this case. In the sparse Hessian approximation case significant error growth occurs when the finite-difference step size varies over the set of finite-difference directions; however, in our current setting there this is not an issue since the "step size" is equal to unity in all cases. First we consider the number of substitutions required to determine any nonzero of J from are chosen using our substitution stategy. There is good news: similar to the sparse Hessian approximation situation [4, 7, 16], the number of dependencies, or substitutions, to resolve a nonzero of J can be bounded above by 1b(m Theorem 5.1 Let OE be a bipartite cyclic p-coloring of G b (J). Then, OE corresponds to a substitution determination of J and each unknown in J is dependent on at most m+n \Gamma 2 unknowns. Moreover, it is possible to order the calculations so that the maximum number of substitutions is less than or equal to 1 First, edges (nonzeros) with one end assigned color "0" can be determined directly: by the definition of bi-coloring there will be no conflict. Second, every pair of positive colors induces a forest (i.e., a collection of trees in G b (J)); therefore, the edges (nonzeros) in the induced forest can be resolved via substitution [4]. Hence, all edges (nonzeros) can be resolved either directly or by a substitution process and the worst case corresponds to a tree with m yielding an upper bound of m+n \Gamma 2 substitutions. However, it is easy to see that the substitution calculations can be ordered to halve the worst-case bound yielding at most 1 substitutions. Next we develop an expression to bound the error in the computed Jacobian. Except for the elements that can be resolved directly, the nonzero elements of the Jacobian matrix can be solved for by considering each subgraph induced by 2 positive colors (directions), one color corresponding to a subset of rows, one color corresponding to a subset of columns. Let us look at the subgraph G p;q induced by colors p (columns), q (rows). Let z q J , R q be the set of rows colored q, and C p be the set of columns colored p, where Let denote the quantities computed via AD. Note that the errors introduced here are only due to the automatic differentiation process and are typically very small. In the solution process an element J ij is determined: z 3 Of course a substitution method does incur the extra cost of performing the substitution calculation. However, this can be done very efficiently and the subsequent cost is usually negligible. depending on whether a column equation (of form Jv) or a row equation (of form w T J) is used. Here, N(r i ) denotes set of neighbours of row i in G p;q , and N(c j ) denotes set of neighbours of column j in G p;q . Assume that J actual denotes the actual Jacobian matrix; hence, J actual J actual or J actual J actual Define an error matrix ij be the difference z depending on the way element J ij was computed. We take into account the effect of rounding errors by letting - ij to be equal to ffl ij plus the contribution from rounding errors due to use of the equation that determines J ij . We can now again, depending on the way J ij is calculated. Moreover, we let ffl max be the constant : Note that ffl max has no contribution from step sizes, unlike results for finite- differencing [4, 16]. Theorem 5.2 If J is obtained by our substitution process then Proof: From equation (5.1), or Let us assume, without loss of generality, that equation (5.2) holds. This implies a bound But the same decomposition can be applied recursively to each E ik , and using Theorem 5.1, the result follows. There are two positive aspects to Theorem 5.2. First, unlike the sparse finite-difference substitution method for Hessian matrices [4, 7, 16], there is no dependence on a variable "step size": in AD the "step size" is effectively uniformly equal to unity. Second, there is no cumulative dependence on nnz(J) but rather just on the matrix dimensions, m+ n. However, there is one unsatisfactory aspect of the bound in Theorem 5.2: the bound is expressed in terms of ffl max , but ffl max is not known to be restricted in magnitude. A similar situation arises in the [4, 7, 16]. Nevertheless, as illustrated in the example discussed below, ffl max is usually modest in practise. We conclude this section with a small experiment where we inspect final accuracies of the computed Jacobian matrices. The test function F is a simple nonlinear function as described in x4.2. In Table 7 "FD1" is the sparse finite difference computation [8] using a fixed stepsize "FD2" refers to the sparse finite difference computation [8] using a variable stepsize: ff is uniformly varied in the range [ ffl]. The column labelled "Rel error" records kERRk 2 , where the nonzeros of ERR correspond to the nonzeros of J : for computed J actual The general trends we observe are the following. First, similar to the results reported in [1] for forward-mode direct determination, the Jacobian matrices determined by our bi-coloring/AD approach are significantly and uniformly more accurate than the finite-difference approximations. This is true for both direct determination and the substitution approach. Second, the direct approach is uniformly more accurate than the substitution method; however, the Jacobian matrices determined via substitution are sufficiently accurate for most purposes, achieving at least 10 digits of accuracy and usually more. Finally, on these problem there is relatively little difference in accuracy between the fixed step method and the variable step method. However, as illustrated in [6], it is easy to construct examples where the variable step approach produces unacceptable accuracy. Direct Substitution FD1 FD2 size Rel error Rel error Rel error Rel error Table 7: Errors (sample nonlinear problem) 6 Concluding remarks We have proposed an effective way to compute a sparse Jacobian matrix, J , using automatic differentiation. Our proposal uses a new graph technique, bi-coloring, to divide the differentiation work between the two modes of automatic differentiation, forward and reverse. The forward mode computes the product JV for a given matrix V ; the reverse mode computes the product W T J for a given matrix W . We have suggested ways to choose thin matrices V; W so that the work to compute the pair (W T J; JV ) is modest and so that the nonzero elements of J can be readily extracted. Our numerical results strongly support the view that bi-coloring/AD is superior to one-sided computations (both AD and FD) with respect to the order of work required. Of course AD approaches offer additional advantages over FD schemes: significantly better accuracy, no need to heuristically determine a step size rule, and the sparsity pattern need not be determined a priori [14]. Implicit in our approach is the assumption that the cost to compute Jv by forward mode AD is equal to the cost of computing w T J by reverse mode AD, where v; w are vectors. This is true in the order of magnitude sense - both computations take time proportional to !(F ) - but the respective constants may differ widely. It may be pragmatic to estimate "weights" w respect to a given AD tool, reflecting the relative costs of forward and reverse modes. It is very easy to introduce weights into algorithm MNCO (x4:3) to heuristically solve a "weighted" problem, is the number of row groups (or colors assigned to the rows), and - 2 is the number of column groups (or colors assigned to the columns). The heuristic MNCO can be changed to address this problem by simply changing the conditional (LB) to: R Different weights produce different allocations of work between forward and reverse modes, skewed to reflect the relative costs. For example, consider a 50-by-50 grid matrix with x5), and let us vary the relative weighting of forward versus reverse mode: w . The results of our weighted bi-coloring approach are given in Table 8. Table 8: Weighted problem results Finally, we note that the bi-coloring ideas can sometimes be used to efficiently determine relatively dense Jacobian matrices provided structural information is known about the function F . For example, suppose is a partially separable function, t, and each component function F i typically depends on only a few components of x. Clearly each Jacobian function J i is sparse while the summation, may or may not be sparse depending on the sparsity patterns. However, if we define a "stacked" function ~ F , ~ then the Jacobian of ~ F is ~ J is sparse and the bi-coloring/AD technique can be applied to ~ J, possibly yielding a dramatic decrease in cost. Specifically, if J is dense (a possibility) then the work to compute J without exploiting structure is n \Delta !(F ) whereas the cost of computing ~ J via bi-coloring/AD is approximately J)) is the minimum number of colors required for a bipartite cyclic coloring of graph G b ( ~ J). Typically, - c (G b ( ~ n. The idea of applying the bi-coloring/AD technique in a structured way is not restricted to partially separable functions [9]. Acknowledgements We are very grateful to Andreas Griewank, his student Jean Utke, and his colleague David Juedes for helping us with the use of ADOL-C. --R New methods to color the vertices of a graph The cyclic coloring problem and estimation of sparse Hessian matrices Structure and efficient Jacobian calculation On the estimation of sparse Jacobian matrices Direct calculation of Newton steps without accumulating Jacobians Computing a sparse Jacobian matrix by rows and columns On the estimation of sparse Hessian matrices --TR --CTR Shahadat Hossain , Trond Steihaug, Sparsity issues in the computation of Jacobian matrices, Proceedings of the 2002 international symposium on Symbolic and algebraic computation, p.123-130, July 07-10, 2002, Lille, France Dominique Villard , Michael B. Monagan, ADrien: an implementation of automatic differentiation in Maple, Proceedings of the 1999 international symposium on Symbolic and algebraic computation, p.221-228, July 28-31, 1999, Vancouver, British Columbia, Canada L. F. Shampine , Robert Ketzscher , Shaun A. Forth, Using AD to solve BVPs in MATLAB, ACM Transactions on Mathematical Software (TOMS), v.31 n.1, p.79-94, March 2005 Shaun A. Forth, An efficient overloaded implementation of forward mode automatic differentiation in MATLAB, ACM Transactions on Mathematical Software (TOMS), v.32 n.2, p.195-222, June 2006 Thomas F. Coleman , Arun Verma, ADMIT-1: automatic differentiation and MATLAB interface toolbox, ACM Transactions on Mathematical Software (TOMS), v.26 n.1, p.150-175, March 2000

sparse finite differencing;NP-complete problems;ADOL-C;automatic differentiation;nonlinear systems of equations;bicoloring;computational differentiation;sparse Jacobian matrices;partition problem;nonlinear least squares;graph coloring

294403

Inertias of Block Band Matrix Completions.

This paper classifies the ranks and inertias of hermitian completion for the partially specified 3 x 3 block band hermitian matrix (also known as a "bordered matrix") P=\pmatrix{A&B&?\cr B^*&C&D\cr ?&D^*&E}. The full set of completion inertias is described in terms of seven linear inequalities involving inertias and ranks of specified submatrices. The minimal completion rank for P is computed.We study the completion inertias of partially specified hermitian block band matrices, using a block generalization of the Dym--Gohberg algorithm. At each inductive step, we use our classification of the possible inertias for hermitian completions of bordered matrices. We show that when all the maximal specified submatrices are invertible, any inertia consistent with Poincar's inequalities is obtainable. These results generalize the nonblock band results of Dancis [SIAM J. Matrix Anal. Appl., 14 (1993), pp. 813--829]. All our results remain valid for real symmetric completions.

Introduction We address the following completion problem: given a partially specified hermitian matrix characterize all the possible inertias In d) of the various hermitian completions H of P: We call this set the "inertial set" or "inertial polygon" of P . The issue of classifying positive definite and semidefinite completions of partial matrices is relevant to various applications involving interpolation, and has been studied thoroughly, e.g. [AHMR], [D5], [GJSW]. Invertible completions have been studied in [DG] and [EGL2], for band patterns, in [L] for general patterns, and are associated with maximum entropy and statistics. For other results concerning ranks and general inertias, see [D1-D6], [EL], [G], [H], [JR1], [CG3]. Following some preliminary material (Sections 2-4), we present in Sections 5 and 6 several contributions to the inertia classification problem. will denote the inertia, that is the number of positive, negative, and zero eigenvalues of a hermitian matrix H. They are also called the positivity, negativity and nullity of H. The main result is: Theorem 1.1 Given the block "bordered" matrix @ Z D EC A are Hermitian (real or complex) matrices of sizes ff \Theta ff, fi \Theta fi and fl \Theta fl, respectively, and B and D are matrices of sizes ff \Theta fi and fi \Theta fl, respectively. Set: In and (- In For given integers n and p, there exist an ff \Theta fl (real or complex, respectively) matrix Z such that In if and only if Figure 1.1. Graph of the inertial polygon for the bordered matrix The block partition is not required to be uniform; rectangular (non-square) blocks are permitted. For this partial matrix P (Z), Theorem 1.1 shows that the inertial set is a (possibly degenerate) convex seven sided lattice polygon as depicted above. The proof of Theorem 1.1 is presented in Section 5, along with a variety of corollaries including a small application to the algebraic matrix Riccati equation. Cain and Sa established the 2x2 case (i.e., [CS]. The result was generalized to an arbitrary number of diagonal blocks by Cain in [C], with further results by Dancis in [D1]. The case with one given diagonal block was proven as Theorem 1 of [S] and as Theorem 1.2 of [D1]. These cases are reviewed in detail in Section 4, and are later used as milestones in computing the inertial polygon of Theorem 1.1. The possible inertias for a bordered matrix missing a single (scalar) entry was catologued by the second author in [D6], mostly using his Extended Poincar'e's Inequalities (3.3). In [JR] the lower bounds in (5.5) and (5.6) were determined for the case when the given principal blocks are invertible (i.e., ff extends to the case of "chordal patterns". In computing the inertial set for Theorem 1.1, we combine four simple elements: (i) Schur complements, (ii) Poincar'e and Extended Poincar'e Inequalities (3.3) as necessary conditions on the inertia, (iii) the technique of "restricted congruence" (presented in Subsection 2.4), including a new formula (2.7), for simplifying a partial hermitian matrix and (iv) elimination of variables in systems of linear inequalities (see Sections 2-3 for details). These techniques enable us to reduce Theorem 1.1 to a combination of the simpler cases presented in Section 4. These four elements of the proof, without (2.7), are commonly used in the matrix literature, and in particular in the completion literature cited above. Staircase hermitian matrices are the mild generalization of block band matrices described as generalized block band matrices in the appendix of [JR2]; they look like a double staircase which is symmetric about and includes the main diagonal. A staircase (or generalized block band) matrix with s steps is a partial hermitian n \Theta n R with precisely s specified hermitian submatrices, fR 1 are defined by @ a a k i 1. The inertia of each of the R i 's is denoted by In R Staircase matrices allow the non-diagonal blocks to be non-square rectangles. Note that R 1 need not overlap R 3 as would be required in a block band matrix. In fact, R 1 need not even but the main diagonal must be contained in the union of the R i 's. Also the definition includes block diagonal matrices. The next theorem shows that a staircase matrix, with all maximal submatrices being invertible, has hermitian completions with all the inertias consistent with Poincare's Inequalities. Theorem 1.2 (An inertial triangle) Given an s-step hermitian staircase m \Theta m matrix R. Suppose that each of the maximal submatrices R 1 s of R is invertible. Then the inertial polygon of R is the triangle The proof of Theorem 1.2 is presented in Section 6, along with several theorems about the possible inertias of hermitian completions of staircase matrices. We will employ the method of Dym and Gohberg [DG], which decomposes the completion process into a succession of simple steps, each of which is a Theorem 1.1 step. These results generalize the (scalar) band hermitian completion results of the second author in [D6]. A related of Johnson and Rodman is restated as Lemma 5.7 herein. We state our results for complex hermitian matrices, but they are all equally valid in the real symmetric case. Preliminaries 2.1. Notation: We shall denote by p; n and p; n the minimal, resp. maximal possible values of the positivity and the negativity of the completions of a given partially specified matrix. Similarly, r and r will denote the minimal and maximal possible values for the rank of completion matrices of a given partially specified matrix. We have the obvious inequality r - p+n; which is generally strict. The determination of the maximal rank r for arbitrary (including non- band) hermitian completion problems is done in [CD]. In fact, it is shown there that the maximal completion rank does not increase if the assumption that the completion is hermitian is dropped; consequently, this rank can be computed explicitly using a result of [CJRW]. The inequality r - p + n is similarly obvious; however, it becomes an equality (i.e. in many cases, including that of Theorem 1.1, see Corollary 5.4, and certain block band matrices, see Theorem 6.2. d) will denote the square matrix I p \Phi \GammaI n \Phi 0 d of inertia (p; n; d) . Sometimes we shall use the triple (-; ffi) to denote the inertia of a given (maximal) specified submatrix and d) to denote the inertia of a hermitian completion of a given partial matrix. Congruence of matrices is denoted by - =. If a square matrix X is written in block form, say ij is of size a i \Theta a j ; we shall describe X as having block sizes (a 2.2. Schur Complements: Let If A is invertible then the Schur complement of A is Haynesworth [H] has shown that H is congruent to A \Phi C \Theta . In particular, In More generally, if H is a k \Theta k block matrix, and T is a subset of f1; :::; kg; let A be the principal submatrix whose block indices are in T : We can move A to the left upper corner by a permutation of the coordinates, and proceed as before, provided A is invertible. This procedure will be referred to as complementation with respect to coordinates T . The block division will always be clear from the context. A similar procedure is available for non-hermitian matrices, yielding the weaker identity for and C We shall refer to this procedure as non-hermitian complementation. 2.3. Canonical Forms: We shall repeatedly use the following terminology: (i) Equivalence canonical form: It is well known that two matrices A and B are equivalent if their exist two invertible matrices S and T such that A = SBT . Every matrix A can be transformed by equivalence to the form I a 0 , where a = rank A. We shall also need the following case of "block equivalence": For every matrix in block form are invertible matrices respectively, such that @ I a A where (ii) Strong congruence canonical form: Every hermitian matrix A of inertia (p; n; d) is congruent to a matrix of the form J(p; n; d): (iii) Weak congruence canonical form: Every hermitian matrix A of rank r is congruent to a matrix of the form A 0 \Phi 0, where A 0 is an invertible r \Theta r matrix. 2.4. Restricted Congruence: If P is a partial matrix, and S is invertible, the matrix can be interpreted as a partial matrix, in the following sense: an entry p 0 ij of P 0 is determined if it is equal to fS HSg ij for every possible completion H of P: We call a restricted congruence if p ij being a specified entry of P implies that p 0 ij is specified in P 0 . There are some similarities between our concept of "restricted congruence" and Ball et al's concept of "lower similarity" in [BGRS]. We will use restricted congruence in two ways: block-diagonal congruence is used to put some (specified or unspecified) blocks of P in canonical form; some row and column operations are used to annihilate blocks in P: In some cases, unspecified blocks may become specified (in fact, annihilated) by congruence. For example, the 1,1 block of I I 0 can be annihilated by the process I !/ Z I I 0 !/ I 0 I I 0 I 0 0C C A can be simplified toB B @ I 0 0C C A as follows:B @ I I \GammaX @ I 0 0C @ I A =B @ I 0 0C Inequalities Necessary conditions. The shape of an inertial polygon for a hermitian completion problem is, to a large extent, determined by a few inequalities relating matrices and submatrices: (i) Rank Interlacing. If A is a k \Theta l rectangular block of the m \Theta m matrix H; we have Inequalities. Let A be a k \Theta k principal submatrix of the m \Theta m hermitian matrix H: Let - i and - i be the ordered eigenvalues of A and H; respectively. The Cauchy interlacing Theorem states that - i - i - i+m\Gammak (see e.g. Theorem 4.3.15 in [HJ]). An equivalent statement are the Poincar'e inequalities: The upper and lower bounds of (3.2) were strengthened in Theorem 1.2 of [D2]; the lower bound was strengthened as follows: (iii) Extended Poincar'e's Inequalities. ([D2]) Given a hermitian matrix in block form, set Then In H - In Inequalities (3.1), (3.2), (3.3) form a set of a priori bounds on completion inertias. In fact, in [CJRW] it is proved that the upper bound in (3.1) is sufficient for the determination of the maximal completion rank in the case of non-hermitian completions, and in [CD] the same is shown in the case of hermitian completions. It turns out that the necessary conditions of type (3.1), (3.2), (3.3) are also sufficient in determining the full inertial polygon in many cases, including the bordered matrix case (Theorem 1.1) and the block diagonal case ([D1]). We shall emphasize cases of sufficiency of these conditions in the text. 4 Towards the 3 \Theta 3 Bordered Case This section paves the way for the analysis of the bordered case, which is carried out in Section 5. The material in this section has independent value, and much of it is well-known. We shall compute the inertial polygon for a 3 \Theta 3 block pattern of the form @ as Lemma 4.8. The special cases A ? and A ? , originally due to Cain and S'a, will also be reviewed, as Lemmas 4.1 and 4.5. We shall also compute the possible inertias of a matrix of the form A+X with inertia limitations on the unknown matrix X as Lemma 4.3. The results of Lemma 4.1 will be used to establish Lemmas 4.3 and 4.8 . The result of Lemma 4.3 will be used to establish Lemma 4.5 which will be used in the proof of Lemma 4.8 which will be used in the proof of Theorem 1.1. The proof of Theorem 1.1 and Lemma 5.10 consist of reduction to the case of (4.1), which itself, is of independent interest. Lemma 4.1 Let be a partially specified hermitian matrix of block sizes (ff; fi). Then the inertial polygon for H is the pentagon that contains all the lattice points (-(H); -(H)) which satisfy the inequalities: Proof. The necessity of (4.2) follows from (3.1) and (3.2). For sufficiency, put H 1 in diagonal form, and complete H to a diagonal matrix. It is easy to show that every inertia in (4.2) is obtained. See also Theorem 1 in [S] and Theorem 1.2 in [D2]. Corollary 4.2 In Lemma 4.1 the extremal values are Moreover, H in Lemma 4.1 admits positive definite, non-negative definite, or invertible completions if and only if H 1 is positive definite, non-negative definite, or rank(H 1 respectively. The following result can be deduced with some effort from Theorem 2 in [S]: Lemma 4.3 Let A and X be m \Theta m hermitian matrices. We consider A to be fixed and X to be a variable matrix with -(X) - a and -(X) - b: Then the possible inertias of are the nonnegative lattice points satisfying the following inequalities: Proof. Necessity is obvious due to Sylvester's inertia principle. For sufficiency, take A to be diagonal, and restrict X to be diagonal as well. It is easy to show that every inertia in (4.3) is obtained. Corollary 4.4 In Lemma 4.3, 0g. Also, A admits positive definite completions if and only if - a and - non-negative definite completions if and only if - a; and invertible completions if and only if ffi - a For the values of n; n in Lemma 4.3, see also [CG3] Lemma 2.2. The following result is due to Cain and S'a. Lemma 4.5 ([CS]) Let be hermitian of block sizes (ff; fl): Then the inertial polygon of H is determined by these inequalities: \Gamma- A generalization of Lemma 4.5 and (4.4) to more than two diagonal blocks can be found in Brian Cain's paper [C]. A short proof of the necessity part of Cain's result was obtained by J. Dancis ([D1] Corollary 11.1 and Lemma 11.2). In the 2 \Theta 2 block case, this proof is given below. Proof of necessity of Inequalities (4.4). be a completion of H. Set F G Noting that rank F we apply the Extended Poincare Inequalities (3.3) to both F and G and we obtain Subtracting this inequality from -(H) yield the right side of (4:4). A symmetric argument produces the left inequality. Proof of Lemma 4.5. We show by reduction to Lemma 4.3 that Inequalities (4.4) describe the inertial polygon. Let be a completion of H. Putting F in weak canonical form, and taking the Schur complement of the new first coordinate, F 0 , we calculate, using (2.2): @ A Hence Putting G \Theta in weak canonical form, we get H @ Putting X 4 in equivalence canonical form, we get @ I r A We set In(G \Theta Removing coordinate 5 and taking the Schur complement of coordinates 1,3, and 4, in H 0 , we calculate using Equations (2.2) and (2.7): We now develop the relevant inequalities involving the dummy variables in (4.6). By Lemma 4.3 the inertial polygon of G \Theta is determined by the inequalities on \Gamma- By Lemma 4.3 again, we compute the inertial polygon for The only restriction on r is the size of X Now (4.4) is obtained from Equations (4.5) - (4.9) by eliminating - - and r: The values of p; p; n; n can easily be computed from Lemma 4.5. The values n and n; were computed in [CS]. Lemma 4.6 [D1] In Lemma 4.5 we have Moreover, there exists a matrix X which simultaneously achieves the minimal possible ranks for F and G namely rank G Proof. The values of p and n follows directly from Lemma 4.5. To show the rest, we may put F and G in strong canonical form: @ I @ I @ Choose the diagonal entries of X 1;1 and X 2;2 by the rule (X 1;1 1. Choose all other entries of X to be zero. We get a completion with the desired minimal ranks. The original result of J. Dancis ([D1] Theorem 1.3) is in fact more general in two respects: first, it extends to more than two block diagonals. Moreover, it is not restricted to minimal ranks: it shows, more generally, that any choice of kernels of a column decomposition as well as any choice of inertia which is consistent with the Extended Poincar'e inequalities, can be obtained: Theorem 4.7 constrained hermitian completion.) [D1] Given hermitian matrices H ii ; In H ii be the block diagonal matrix of size Choose a subspace K i ae ker H ii such that dim ker H ii \Gamma dim K i - s: Then an integer triple (-; ffi) satisfying the equality - is the inertia of a hermitian completion H of S with column block structure: each M i is an n \Theta n i matrix; and ker M if and only if (-; ffi) satisfies the inequalities (The notation here is different than the one used in [D1]: the r i and \Delta i here correspond to - The next lemma is the main result of this section; it combines Lemmas 4.1 and 4.5. Lemma 4.8 Let P be a partial matrix of the formB B @ A and of block sizes (ff; fl; ffl): Then the inertial polygon of P consists of the lattice points determined by these inequalities: The sufficiency proof of Inequalities (4.12) is the same as for Inequalities (4.4). Proof. Every completion H of P has the form F G was computed in Lemma 4.5. By Lemma 4.1, In(H 1 ) and In(H) are connected by and eliminating In(H 1 ) from Inequalities (4.4) and (4.13), and using the identities we get Inequalities (4.12). 5 Inertias of Block Bordered Matrices In this section we establish Theorem 1.1 using the results stated in Sections 3 and 4. The material in this section is new. The scalar case was classified in [D3]. Other special cases of Theorem 1.1 occur in [D1], [L],[G], and [CG3]. Subsections 5.3-5.5 contain additional results and corollaries of of Theorem 1.1 concerning minimal rank completions of various types, and the case where the two maximal specified hermitian submatrices R 1 and R 2 of P (Z) of Theorem 1.1 are invertible. In Subsection 5.6 we present a small application to the algebraic matrix Riccati equation A+AZ +ZB +ZCZ 0: a criterion for solvability and a characterization of the possible inertias of the solution matrix Z (which need not be hermitian). 5.1. Internal Relations for Bordered Matrices. For the the bordered matrix P (Z) of Theorem 1.1, we note that R ; and R are the maximal specified hermitian submatrices of P (Z) and the maximal specified non-hermitian submatrix of P (Z). Observation 5.1 (Internal relations for a bordered matrix) With the notation of Theorem 1.1, we define and Then Proof. The inequality of (5.1) follows from rank considerations. The equality, as well as (5.4), follows from the definitions. Applying the Extended Poincar'e's Inequalities to C as a submatrix of R 1 or R 2 provides (5.2). (5.3) also follows from the Extended Poincar'e's Inequalities. 5.2. Proof of Theorem 1.1. Before proving the theorem, let us comment on the necessity and minimality of its conditions: Minimality: The inertial polygon for a simple matrix, with all seven edges present, is illustrated in Fig. 5.1. The matrix chosen was of block sizes (6; 1; and B; C and D are zero matrices of appropriate order. This results in r and the three \Delta's being zero. This example shows that the set of seven inequalities defining the inertial diamond is not redundant. Figure 5.1. A seven-sided inertial polygon. Necessity: Necessity of each one of the seven inequalities can be easily demonstrated: inequality (1.4) follows from (3.1). The upper bounds in Inequalities (1.1) - (1.2) are a consequence of (3.2). The lower bounds in (1.1) and (1.2) are just the Extended poincar'e's inequalities (3.3). It remains to derive (1.3). 5.2 The Extended Poincar'e's Inequalities (3.3) imply (1.3). Proof. We define the partial matrices and their completions Z D E Let Using the Extended Poincar'e's Inequalities(3.3) twice, we have: Substitute for one of the p's on the left hand side Using the identities - We note that But this translates into: or equivalently Finally, (5.8) and (5.9) establish (1.3). We will establish Theorem 1.1 by using Schur complements (Equation (2.2)) and row and column operations (Equation (2.7)) and the other forms presented in Section 2, repeatedly, in order to reduce Theorem 1.1 to Lemma 4.8. Proof of Theorem 1.1 We begin by putting C in weak canonical form: with block sizes (ff; rank C; Taking the Schur complement of C 0 in P (Z) as in yields In In - In where @ Y D" GC C Next we put [B" ; D"] in the canonical form (2.4). Using (5.10), we obtain the matrix @ F F I 22 X A The block sizes here are are blocks of Y , the F ij 's and G ij 's are conforming blocks of F and G. The new block sizes are related to ff; fi; fl via Next we use restricted congruence, see Section 2. By row and column operations based on and H , we may assume without loss of generality, that are all zero. This modifies the matrix - without changing its inertia, to @ I \Delta 00 22 X A We may discard row 7 and column 7, which are all zero. Next we complement H 0 with respect to the block [1;2;4;5;6;10] =B @ I \Delta 00 A The Schur complement turns out to be @ @ 22 @ I @ A @ 22 @ with block sizes (ff . By the Schur complement inertia formula (2.2), we get In In H 0 In In H 00 (The fi 0 accounts for removing coordinate 7 from H 0 ). Thus we have In In In We will use Equation (5.17) and Lemma 4.8 to calculate In P (Z); to find In H 00 , we must first calculate In F 33 and In G 11 : 5.3 In F In In C 0 In G In In Proof. By restricting Equations (5.10)-(5.14) to the upper left corner R 1 , we may take the Schur complement of C 0 as a submatrix of R 1 ; this yields: In @ In In F B" where F is as in Equation (5.11). Using elimination, Equation (2.7), we note that In F B" @ I A In F 33 where F 33 is as in (5.12). Solving for In F 33 in Equations (5.19) and (5.20) yields the first part of (5.18). A similar argument holds for the second part. We proceed with the proof of Theorem 1.1. Applying the size identities and Lemma 4.8 (with its submatrices F and G corresponding to F 33 and G 11 here), we obtain these inequalities for In H 00 : Inequalities are obtained by plugging Inequalities (5.21) and Equation (5.18) into Equation (5.17) and then using Equations (5.1), (5.4) and (5.13) to eliminate all the intermediary inertias. 5.3. Extremal inertia values and inertia preserving completions In this subsection and the next, we use the geometry of the inertial polygon as the basis (i) for establishing a minimum rank completion for the bordered matrix P of Theorem 1.1 (Corollary 5.5); and (ii) for showing that, assuming invertibility of R 1 and R 2 , all the inertias which are consistent with Poincare's Inequalities can be obtained by completion (Corollary 5.8). In Section 6 we will use these results as building blocks for our proofs of completion theorems for "staircase" matrices. First we show, under the notation of Theorem 1.1, that there is always a completion whose positivity and negativity are the minimal ones allowed by the Extended Poincar'e's Inequalities. This implies that the minimal rank is Corollary 5.4 (Minimal rank completions.) With the notation of Theorem 1.1, there exists a hermitian completion Proof. We argue by inspection on Figure 5.1. If the vertex not in the inertial polygon for P (z), this vertex must be cut off by one of the extremal lines defining Inequalities (1.2 - 1.4). Since these inequalities are consistent, it is clear that (p must satisfy (1.2) and (1.4). It remains to check the two inequalities (1.3). We have to consider four possible choices for p 0 and in the notation of Theorem 1.1: Suppose that Rank considerations imply that Rank C Equation (5.1) implies that - This and Equation (5.22) imply that proving the right inequality in (1.3). Interchanging the roles of the - 0 s and the - 0 s establishes the left inequality. (II) A similar proof applies to the case (p 0 then a combination of Equations (5.1),(5.3) and (5.4) yields rank R 1 algebra we get - This and (5.23) implies that proving the right inequality in (1.3). Interchanging the roles of the - 0 s and the - 0 s establishes the left inequality. (IV) A similar proof applies to the case Remark. This Corollary 5.4 implies that the minimal rank of the set of hermitian completions of Constantinescu and A. Gheondea presented in [CG3] another formula for n. Of particular interest is the case where the minimal rank solutions inherit their inertia values and n from the specified blocks R 1 and R 2 , i.e. We call such completions inertia preserving. Note that (5.24) does not guarantee that the minimal completion rank is g: The following simple result will play a major role in finding inertia preserving completions for block band matrices in Section Corollary 5.5 (Inertia preserving completions.) Assume the notation of Theorem 1.1. Suppose that P satisfies an "equality of ranks" condition: Then P admits inertia preserving completions. Indeed, under condition (5.25), the formulas of the inertial polygon simplify, since we have In particular, (5.24) holds. Condition (5.24) and Corollary 5.5 are implied by the stronger condition In the notation of Observation 5.1, Equation 5.26 is equivalent to any of the following: This condition is satisfied if e.g. C is invertible. Condition (5.25) is not necessary for the existence of inertia preserving completions. As an example, consider the partial matrix @ A simple argument, using the Extended Poincar'e's Inequalities(3.3), shows that if (5.25) is not satisfied, then the inertia preserving completion must inherit both p and n, from the same block. Hence In the above example, 5.4. The width of the inertial set We define the width w of the inertial polygon to be the maximal value of all pairs of points (p; n) and (p belonging to the inertial polygon. See Figure 1.1. This width equals the sum of the lengths of the two perpendicular sides of the inertial polygon. It is clear that Inequality (1.3) puts a limitation on the width; namely, w cannot exceed the modulus of the difference of the right and left hand sides in this inequality: (it can be shown directly that this value is always non-negative). The sides, with slope of minus 1, come from Inequality (1.3), which may be rewritten as In this way is related to the width of the inertial polygon. The width of the inertial polygon tends to increase as we increase the ranks of R 1 and R 2 : In this section we study the two extreme cases. The "slim" case is when Rank R Here the polygon degenerates to a segment with a degree inclination. The "fat" case occurs under the maximal rank condition det R 1 det R 2 6= 0; here the polygon extends to maximum capacity, and fills a triangle (Corollary 5.8 ). We start with the "slim" case. Corollary 5.6 Given the notation of Theorem 1.1. Suppose that Then the inertial polygon coincides with the segment Moreover, the minimal rank completion is unique. Proof. Condition (5.30) together with the Poincar'e inequalities imply that C; R 1 and R 2 all have the same inertia (-; ) (with possibly different nullities). The same condition also implies (5.26), hence (5.25), and Corollary 5.5 can be used. We conclude that That condition (5.31) implies zero width, is clear from (5.27). The rest which restricts the polygon to a line segment of the form (- K: The value follows from Theorem 1.1 with some algebra. It can also be deduced from the maximal rank considerations in [CD]. To prove uniqueness of the minimal rank completion, we note that (5.30) forces the factorizations I I can check that the completion rank is rank C the unique solution requires that Z Observation 5.6 represents the extreme case of a "slim" inertial set. We now turn to examine the other extreme case of a "fat" inertial set. Under the assumption that R 1 and R 2 are invertible matrices, four inequalities among (1.1-1.4) are redundant, and the inertial polygon becomes a triangle, admitting any inertia compatible with the Poincar'e inequalities and the size limitation. First we quote the following known result about matrices with chordal graphs. Chordality is discussed in [GJSW], [JR1] and [JR2], and it suffices to say that block bordered 3 \Theta 3 patterns (and in fact the general staircase patterns of Section have chordal graphs. Lemma 5.7 . (Corollary 6 of [JR1]) In any hermitian partial matrix P of size m \Theta m whose pattern has a chordal graph, and all its maximal hermitian specified submatrices are invertible, the points together with all the lattice points on the straight line segment connecting them, belong to the inertial set of P: In the bordered case we can say more: Corollary 5.8 Assume the notation of Theorem 1.1. Suppose that R 1 and R 2 are invertible. Then the inertial polygon is the triangle whose vertices are In other words, every inertia consistent with the Poincar'e inequalities Proof. Let T be the triangle defined by the above three inequalities. Let D be the inertial polygon. It is easy to check that v 0 are the three vertices of T : Since the Poincar'e inequalities are a subset of (1.1) - (1.4), we get the inclusion D ae T : Note that in (1.4) our hypothesis implies that On the other hand, v By convexity, we conclude that T ae D: 5.5. Simultaneous rank minimization We now strengthen the minimal rank result obtained in the last subsection (Corollary 5.4). Consider the partial matrices N 0 and N 00 of Equation 5.5. We wish to find a matrix Z which will simultaneously induce minimal rank completions in N 0 and N 00 as well as in the full bordered matrix P . Before we tackle the general case, let us make the simplifying assumption (5.26), for which a slightly stronger result is available. This very simple special case also serves as an outline and motivation for the general case. Also readers who are only interested in Theorems 6.2 and 1.2 and Corollary 6.4 but not in Theorem 6.7 may read the proof of Lemma 5.9 and skip the calculations of Lemma 5.10. Lemma 5.9 Assume, along with the notation of Theorem 1.1, that Then there exists a matrix Z 0 satisfying simultaneously the inertia preserving condition and (using the notation of Equations 5.5 and 5.6) the two minimal rank conditions Such a completion also satisfies the kernel condition Ker Proof. Since Equation (5.26) implies Corollary 5.5, P admits inertia preserving completions, (5.32). The fact that rank(R 1 ) and Rank(R 2 ) are the minimal completion ranks for N 0 and N 00 is obvious. To prove that conditions (5.32-5.33) are attainable simultaneously, we re-examine the proof of Theorem 1.1, and reduce the situation to Lemma 4.6, where a positive answer is available. We assume the rank condition (5.26), which implies - follow the proof of Theorem 1.1. The matrices B 00 and D 00 in (5.10) turn out to be zero: @ A of block sizes (ff; rank C; ffi(C); fl) . Now removing the zero row and column and then taking the Schur complement with respect to C 0 yields Y G We get therefore Y G Lemma 4.8 shows that a matrix exists for which - G are simultaneously minimum rank completions. Choosing Z we see from (5.34) that Z 0 minimizes all the three ranks involved. It remains to show the kernel condition. First, we observe, for all Z, that Ker P (Z) oe Ker (N 0 (Z) \Phi I) and Ker(N 0 (Z)) ae Ker(R 1 For Z 0 just obtained, we actually have (5.33), hence the second containment must be equality: Now the first containment becomes Ker Similarly Ker P (Z 0 ) oe Ker (I \Phi R 2 Ker In the general situation, when the simplifying assumption rank is not assumed, a simultaneous minimal rank solution still exists, but it is not necessarily an inertia preserving solution, and the additional kernel condition cannot be guaranteed. Lemma 5.10 simultaneous minimal rank completion lemma) With the notation of Theorem 1.1, the minimal completion ranks for are, respectively, Moreover, there exists a matrix Z 0 which produces these ranks simultaneously. Proof. Assume the notation of Theorem 1.1 and Observation 5.1. First we verify the expressions for the minimal ranks involved. Corollary 5.4 implies that matrices. Using our definitions of \Delta 0 and \Delta 00 , the identities r(N 0 are obvious. Having computed the minimum completion ranks for these 3 matrices, we now demonstrate that the three minimum ranks can be achieved simultaneously. Our plan is to perform all the steps of the proof of Theorem 1.1 simultaneously on the three matrices involved. We call a step permissible if a completion exists which preserves the three minimal ranks. As will be seen, not all steps are permissible, and some modification will be necessary. The reduction of H to - H in (5.12) is permissible, since C is a common block in all three matrices. Besides - H; this reduction applied to N 0 and N 00 yields @ F F I A @ I A Using (2.3), these operations preserve ranks: Our aim now is to minimize simultaneously rank - H in (5.12) and rank - The passage from - H in (5.12) to H 0 in (5.14) is also permissible, and may be followed by a similar passage from - H i to new matrices H 0 all the F; G; X entries located in first and last block rows and columns in - H i are made zero. We also discard zero rows and columns in these matrices (the seventh block coordinate in H 0 ). Permissibility is violated in the passage from H 0 to H 00 in (5.15). More precisely, complementation of H 0 with respect to coordinates 1,4,6,10 is permissible; unfortunately, symmetric complementation with respect to coordinates 2 and 5 is not permissible, since these coordinates are not present in both H 0 Consequently, the proof of Theorem 1.1 has to be modified: we perform on H 0 non-hermitian complementation (2.3) with respect to block rows 1,2,5,6 and block columns 4,5,9,10, i.e. with respect to the matrix [1;2;5;6][4;5;9;10] =B @ I \Delta 00 A . The Schur complement of H 0 with respect to H 0 [1;2;5;6][4;5;9;10] isB @ I \Delta 00 22 X A \GammaB @ A @ I \Delta 00 A @ A @ I 22 X A This matrix is of the general formB @ I A where the W i 's are unspecified. Indeed, it is easy to see that any arbitrary choice of the W i 's can be achieved by appropriate choice of the X i 's. The respective Schur complements of H 0 1 and 2 with respect to H 0 [1;2;5;6][4;5;9;10] turn out to be ~ I @ A Taking zero is obviously a minimal rank choice for all three matrices. We have reduced the original problem to the simpler problem of simultaneously minimizing the ranks of the three matrices F G 11 Reduction to Lemma 4.8 is completed. In Section 6 we shall use the following weakened form of Lemma 5.10, which has better propagation properties. Corollary 5.11 (Propogation of internal inequalities of Bordered Matrices) Given the notation of Theorem 1.1 and Observation 5.1, then there exists a matrix Z 1 such that Corollary 5.11 follows directly from Lemma 5.10 and Observation 5.1. 5.6. Solvability of the Ricatti equation. We end this section with a small contribution connected to the theory of Lyapunov and Riccati equations. Lemma 5.12 Given matrices A; B, and C of sizes ff \Theta ff; al \Theta fi; fi \Theta fi; respectively, with A and C hermitian, define Then the possible inertias of matrices of the form P arbitrary Z, form the septagon Proof. This is an easy corollary of Theorem 1.1, using complementation on the last two block coordinates of the bordered matrixB B @ Z I 0C C In [CG3] the values of n and n were determined for this case. Corollary 5.13 The Riccati equation A solvable if and only if These results apply also for the Lyapunov or Stein equations: simply assume that C or B is a zero matrix. We emphasize, however, that in the classical context of these equations Z is assumed hermitian (at least), and then it is not clear whether this puts additional restrictions on the set of inertias. 6 Some Completion Results for General Band Matrice In this section, we consider hermitian matrices with general block band or "staircase structure", where again the blocks may vary in size. We follow the method of Dym and Gohberg [DG], which decomposes the completion process into a succession of simple steps, each involving the completion of one bordered submatrix. Combining this procedure with the results of Section 5, we are able to draw several interesting conclusions: In Subsection 6.3 we identify certain classes of staircase hermitian matrices which admit inertia preserving completions; these are completions which inherit their inertia values p and n from (possibly two different) specified blocks of the given partial matrix P . Such completions are obviously minimal rank completions (See Theorem 6.2 and Corollaries 6.3 and 6.9). These results generalize the (scalar) band hermitian completion results of the second author in [D6]. In Subsection 6.3 we consider block band or staircase matrices for which all maximal specified hermitian submatrices are invertible. We show that such matrices admit all the possible completion inertias consistent with Poincar'e's inequalities (See Theorem 1.2) which includes inertia preserving completions. These results generalize the (scalar) band hermitian completion results of the second author in [D6]. (III) Not every partial matrix admits inertia preserving completions. In Subsection 6.4 we establish modest upper bounds on the minimal posssible rank for hermitian completions of staircase hermitian matrices. 6.1. The Staircase Matrix Notation In dealing with staircase partial matrices, we shall adhere to the following notation and observations, which shall collectively be referred to as the Staircase Notation. We recognize the bordered matrix of Theorem 1.1 in each pair (R i and R i+1 ) of successive maximal hermitian submatrices. We therefore define the s bordered partial submatrices P i to be: @ where these specified submatrices A of R are As in the notation of Theorem 1.1, we observe that the R i are the specified block submatrices: and that each C i is the overlap of R i and R i+1 . The submatrices R 1 ; C and R 2 of the notation of Theorem 1.1 correspond to the submatrices R of each bordered submatrix P i . We shall denote by P i (Z i ) the completions of P i , using Z i in the right upper block of We denote the inertias of P i (Z i ) by (- 0 In dealing with the i'th bordered submatrix P i ; the incremental ranks d defined in the notation of Theorem 1.1 and Observation 5.1will be distinguished by the subscript i. 6.2. The Diagonal Completion Formalism diagonal [partial] completion R+ of an s-step staircase partial matrix R is an staircase partial hermitian matrix obtained by completing all the s bordered matrices P i of R. This entails the addition of a matched pair of whole block diagonals alongside the specified band. In a diagonal completion the different bordered completions P i (Z i ) are independent of each other, that is the Z i do not overlap. (II) A standard procedure for completing a staircase matrix is by a succession of diagonal completions. Let F denote a full hermitian completion of R. We may obtain F via a chain of s staircase partial matrices, Each matrix in (6.1) is obtained from its precursor via diagonal completion, and its staircase pattern is reduced by one step. We shall distinguish between the relevant submatrices of each matrix in this chain by attaching to them the appropriate number of (subscript) plus signs. (III) The N st matrix (1 - in the above chain is called an N-diagonal [partial] completion of R. We will identify certain submatrices of R with their counterparts in R+ : For example, the completed bordered matrices P i (Z i ) of R will be identified with the matrices R+i of R+ : In addition, the maximal submatrices R i of R will be identified with the submatrices C+i of R+ : The ordering of these matrices will always be from top left to bottom right. This technique of completing a hermitian scalar band matrix by adding successive pairs of diagonals was developed by Dym and Gohberg in [DG]. This technique of completing a hermitian scalar band matrix was later used in many papers, including [D5], [D6], [DG], [EGL1], [EGL1] and [EL]. We shall use the more general staircase or general block band approach of the appendix of [JR2]. In applying the bordered case (for example, Theorem 1.1) to sections of a band matrix or the more general staircase matrices, the key concept is propagation. Those properties which survive a single (Theorem 1.1) completion step, will by induction, survive the full completion process. Our results in this section are all based on properties which propagate. 6.3. Inertia preserving completions We call an N-step completion R 0 of R inertia preserving if it satisfies the equations: In particular, if F is a fully specified completion then it is inertia preserving if Such a completion is necessarily a minimal rank completion. Not every partial staircase matrix admits an inertia preserving completion. In fact, Example 6.10 will present an infinite sequence of partial staircase matrices whose maximal specified submatrices all have rank 2 but all the full hermitian completions are invertible. Lemma 6.1 (Propagation of inertia preservation) Given an s-step hermitian staircase matrix R. Using the notation of Subsection 6.1, suppose that the blocks of R satisfy these propagation equations: Rank (B for each (using the notation of Subsection 6.2): (i) There exists a one step inertia preserving completion R+ of R for which Rank (B for each (ii) This completion R+ satisfies the kernel condition Ker R+i oe Ker (R i \Phi I) Proof. The proof of (i) is a straightforward application of Lemma 5.9 repeated s times. (6.4) implies that each P i fulfills the hypothesis of that lemma. Therefore each P i admits a completion simultaneously using the notation of Equations 5.5 and 5.6. For the condition rank N into For the condition rank N Rank (B Together this is precisely (6.5). The condition -(P of Lemma 5.7 applied to each Part (ii) follows from part (i), (3.3) and Lemma 5.9. We observe that (6.4) is a propagation condition: Lemma 6.1 shows that this condition can be made to survive a single diagonal completion. Repeating Lemma 5.12 s times along the chain (6.1), we get Theorem 6.2 (Inertia preserving completions) Given an s-step hermitian staircase matrix R fulfilling the propagation Equation (6.4): rank (B Then R admits an inertia preserving fully specified completion F for which Moreover, for this completion F , Ker F contains all the appropriate kernels of the form Ker (I \Phi Theorem 6.2 and its proof are largely a block generalization of Dancis' proof in [D6]. The inequalities show that the expressions in the theorem are lower bounds for Theorem 6.2 shows that they are achieved. Corollary 6.3 If all the matrices, C i and R i in a staircase matrix R, have the same rank r, then there exists a hermitian completion F with rank r. Proof. The condition that all the C i and R i matrices have the same rank implies (6.4) and hence Theorem 6.2 is applicable. This corollary was established for (hermitian and nonhermitian) completions of hermitian and nonhermitian, resp., band matrices in [EL]. An important special case where Theorem 6.2 is applicable is when all the C i submatrices of R are invertible. Corollary 6.4 (Inertia preserving completions) Given an s-step staircase hermitian matrix R. Suppose that the s submatrices C i of R are all invertible. Then R admits an inertia preserving completion F whose kernel contains all the appropriate kernels of the form Ker (I \Phi R i \Phi I): Proof Since all the C i are invertible, the propogation Equation (6.4) holds and Theorem 6.2 applies. 6.4. Incremental bounds on inertia growth In general, even assuming a minimal rank completion in each step, the ranks of the matrices in (6.1) may increase. At present, we cannot compute the minimal rank for the completions of a general staircase matrix or even for a general scalar band matrix. The reason is lack of propagation: the inertias of P i (Z i ) do not depend exclusively on the inertias of R i and C i ; as is evident from Theorem 1.1. However, use of Corollary 5.11 will enable us to obtain an upper bound on the inertia increase. In the next observation and lemma, the B+i ; C+i and D+i matrices are the B matrices of a diagonal completion, R+ ; the d+i will be which is consistent with our general notation. Observation 6.5 Let R be a staircase matrix, together with the notation of Subsections 6.1 and 6.2. Suppose that R+ is a diagonal completion of R for which all bordered completions satisfy the simultaneous minimal rank completion Lemma 5.10. Then the incremental ranks of the bordered submatrices of R and of R+ are related via i and d 00 Proof. We have the following connections between R-related and R+ -related objects: From Lemma 5.10 we note that Combining these equations yields d 0 . The other equation may be similarly observed. Lemma 6.6 (Propagation of incremental ranks) Let R be an s-step hermitian staircase ma- trix. We use the notation which is consistent with our bordered and band matrix notation. Set - d+i g: Then there exists a diagonal completion R+ of R for which Proof. The proof is a straightforward application of Corollary 5.11. The definition of - d and implies that - Therefore for each P i , there exists a completion P i (Z i ) which achieves simultaneously and The last inequalities (6.9) translate to d 0 d and d 00 d; which repeated over all i implies that d: The former inequalities (6.8) prove the rest of (6.7). We observe that Inequalities (??) combine to form a propagation condition: Lemma 6.6 shows that these inequalities may be transferred from a staircase matrix to a diagonal completion. Repeating Lemma 6.6 untill R is fully completed, we get: Theorem 6.7 (Incremental bounds on inertia growth) Given an s-step hermitian staircase matrix R (together with the notation of Subsection 6.1) then there exists a hermitian completion F of R whose inertia (p; n; d) satisfies Corollary 6.8 Given a hermitian staircase matrix R (together with the notation of Subsection 6.1). Suppose there is an integer t such that then there exists a hermitian completion F of R such that st st Proof. The hypotheses and (5.3) provide: rank C d i are integers, this inequality becomes - t. In this way, one sees that t - maxf - and the theorem is applicable. The case is of particular interest: Corollary 6.9 Given a hermitian staircase matrix R (together with the notation of Subsection 6.1). Suppose that then there exists an inertia preserving hermitian completion F of R . That Theorem 6.7 is best possible without additional hypotheses is demonstrated by the next example: Example 6.10 Consider the matrix I where U is an strictly upper triangular matrix. We consider R(U) as a partial s-step band matrix, in which U is unspecified. In the notation of Theorem 6.7, all - Thus, by this theorem, we expect to find a completion with However, since det hence (6.10) can only be satisfied with equality. In fact, using the extended Poincar'e inequalities, we see that every completion R(U) satisfies (6.10) with equality. 6.5. Proof of Theorem 1.2. - a staircase of invertible maximal hermitian submatrices Throughout this subsection we shall assume that R is an s-step staircase partial matrix with all maximal submatrices R i invertible. In this case, all inertias compatible with Poincares Inequalities are achievable with a hermitian completion. First we present the minimal-rank inertia-preserving case as the next lemma. Lemma 6.11 (An inertia preserving lemma) Given an s-step staircase hermitian matrix R (together with the notation of Subsection 6.1). Suppose that each of the maximal submatrices s of R is an invertible matrix. Then there is a hermitian completion F of R such that and such that Ker F contains Ker R 1 +Ker R 2 Proof. We use Corollary 5.13 s times as we construct an inertia preserving diagonal completion R+ of R. Then R+ will satisfy the hypotheses of Corollary 6.4, which will produce the desired hermitian completion F with -(F Proof of Theorem 1.2. We will use Corollary 5.8 repeatedly to construct successive diagonal completions with invertible maximal matrices and increasing inertias. Construction of the (first) diagonal completion (R+ ). Case 1. If the size of P then Corollary 5.8 is used to choose a matrix Z i such that is an invertible matrix with Case 2. If the size of P then Corollary 5.8 is used to choose a matrix Z i such that In Depending only on its size, P i (Z i ) may be an invertible or a non-invertible matrix. In both cases, the new fC+i g are the previous maximal submatrices, fR i g, and hence all the new fC+i g of R+ are invertible matrices. If Case 2 occurred, at least once, then the desired positivity and negativity has been achieved. Then Corollary 6.4 applied to R+ will provide the desired full completion F , with In If no Case 2 has occured, only Case 1, then all the maximal submatrices of R+ are invertible. In this manner, one constructs a number of successive diagonal completions until Case 2 is used. With each successive diagonal completion, the values of the positivity and negativity grow. At some point, at least one of the new -(R i ) and one of the new -(R j ) (for the latest successive diagonal completion reach the desired values p and n . Then This will occur when Case 2 is used, possibly sooner. With the possible exception of the current diagonal completion, all the maximal specified submatrices of the various successive diagonal completions were invertible (since only Case 1 was used). Therefore all the C i of the current diagonal completion were the invertible maximal submatrices of the previous diagonal completion. Therefore Corollary 6.4 is applicable and it completes the proof. --R Positive definite matrices with a given sparsity pattern. On the eigenvalues of matrices with given upper triangular part Determinantal formulae for matrix completions associated with chordal graphs. The inertia of a Hermitian matrix having prescribed diagonal blocks. Maximal rank Hermitian completions of partially specified Hermitian matrices. Minimal signature in lifting of operators I Minimal signature in lifting of operators II The negative signature of some Hermitian matrices Ranks of completions of partial matrices. The possible inertias for a Hermitian matrix and its principal submatrices. On the inertias of symmetric matrices and bounded self-adjoint operators Several consequences of an inertia theorem. Positive semidefinite completions of partial Hermitian matrices. Choosing the inertias for completions of certain partially specified matrices Ranks and Inertias of Hermitian Toeplitz matrices Report Extensions of band matrices with band inverses. maximum entropy and the permanence principle. Invertible self adjoint extensions of band matrices and their entropy. Completing Hermitian partial matrices with minimal negative signature. Determination of the ii On the inertia of some classes of partitioned matrices. Matrix Analysis I Inertia possibilities for completions of partial Hermitian matrices. Chordal inheritance principles and positive definite completions of partial matrices over function rings chordal graphs and matrix completions. On the matrix equation AX --TR

matrices;minimal rank;inertia;hermitian;completion

295659

Finitary fairness.

Fairness is a mathematical abstraction: in a multiprogramming environment, fairness abstracts the details of admissible (fair) schedulers; in a distributed environment, fairness abstracts the relative speeds of processors. We argue that the standard definition of fairness often is unnecessarily weak and can be replaced by the stronger, yet still abstract, notion of finitary fairness. While standard weak fairness requires that no enabled transition is postponed forever, finitary weak fairness requires that for every computation of a system there is an unknown bound k such that no enabled transition is postponed more than k consecutive times. In general, the finitary restriction fin(F) of any given fairness requirement Fis the union of all &ohgr;-regular safety properties contained in F. The adequacy of the proposed abstraction is shown in two ways. Suppose we prove a program property under the assumption of finitary fairness. In a multiprogramming environment, the program then satisfies the property for all fair finite-state schedulers. In a distributed environment, the program then satisfies the property for all choices of lower and upper bounds on the speeds (or timings) of processors. The benefits of finitary fairness are twofold. First, the proof rules for verifying liveness properties of concurrent programs are simplified: well-founded induction over the natural numbers is adequate to prove termination under finitary fairness. Second, the fundamental problem of consensus in a faulty asynchronous distributed environment can be solved assuming finitary fairness.

Introduction Interleaving semantics provides an elegant and abstract way of modeling concurrent computation. In this approach, a computation of a concurrent system is obtained by letting, at each step, one of the enabled processes execute an atomic instruction. If all interleaving computations of a system satisfy a property, then the property holds for all implementations of the program independent of whether the tasks are multiprogrammed on the same processor and which scheduling policy is used, or whether the system is distributed and what the speeds of different processors are. Furthermore, the interleaving model is very simple as it reduces concurrency to nondeterminism. A preliminary version of this paper appears in Proceedings of the Ninth IEEE Symposium on Logic in Computer Science, pp. 52-61, 1994. y On leave from Bell Laboratories, Lucent Technologies. z Supported in part by the ONR YIP award N00014-95-1-0520, by the NSF CAREER award CCR-9501708, by the NSF grant CCR-9504469, by the AFOSR contract F49620-93-1-0056, by the ARPA grant NAG2-892, and by the contract 95-DC-324A. The interleaving abstraction is adequate for proving safety properties of systems (a safety property is of the form "something bad never happens," for example, mutual exclusion). However, it is usually not suitable to prove guarantee properties (a guarantee property is of the form "something good will eventually happen," for example, termination) or more general liveness properties. The traditional approach to establishing guarantee properties is to require that all fair computations, instead of all computations, satisfy the property. Intuitively, fairness means that no individual process is ignored forever. Since all reasonable implementations of the system, whether in multi-programming or in multiprocessing, are expected to be fair, if we prove that a program satisfies a property under the assumption of fairness, it follows that the property holds for all possible implementations of the program. While the theory of specification and verification using different forms of fairness is well understood (see, for example, [LPS82, Fra86, MP91]), fairness has two major drawbacks. First, the mathematical treatment of fairness, both in verification and in semantics, is complicated and requires higher ordinals. Second, fairness is too weak to yield a suitable model for fault-tolerant distributed computing. This is illustrated by the celebrated result of Fischer, Lynch, and Paterson that, under the standard fairness assumption, processes cannot reach agreement in an asynchronous distributed system even if one process fails. We quote from their paper [FLP85]: These results do not show that such problems [distributed consensus] cannot be solved in practice; rather, they point out the need for more refined models of distributed computing that better reflect realistic assumptions about processor and communication timings. We propose one such "more refined" model by introducing the notion of finitary fairness. We argue that finitary fairness (1) is sufficiently abstract to capture all possible implementations, both in the context of multiprogramming and in the context of distributed computing, and (2) does not suffer from either of the two aforementioned disadvantages associated with the standard notion of fairness. Justification of finitary fairness A fairness requirement is specified as a subset F of the set of all possible ways of scheduling different processes of a program. Let us first consider a multiprogramming environment, where all tasks are scheduled on a single processor. A scheduler that meets a given fairness requirement F is a program whose language (i.e., set of computations) is contained in F . The language of any program is a safety property (i.e., it is closed under limits). Furthermore, if the scheduler is finite-state, then its language is !-regular. Thus, to capture all finite-state schedulers that implement F , it suffices to consider the (countable) union of all !-regular safety properties that are contained in F . There are several popular definitions of F , such as strong fairness, weak fairness, etc. [LPS82, Fra86]. For every choice of F , we obtain its finitary version fin(F ) as the union of all !-regular safety properties contained in F . In the case of weak fairness F , we show that the finitary version fin(F ) is particularly intuitive: while F prohibits a schedule if it postpones a task forever, fin(F ) also prohibits a schedule if there is no bound on how many consecutive times a task is postponed. In general, a fairness requirement F is an !-regular liveness property [AFK88]. We show that the finitary version fin(F ), then, is still live, but not !-regular. Now let us consider a distributed environment, where all tasks are executed concurrently on different processors. Here, finitary fairness corresponds to the assumption that the execution speeds of all processors stay within certain unknown, but fixed, bounds. Formally, a distributed system can be modeled as a transition system that imposes lower and upper time bounds on the transitions [HMP94]. We show that a timed transition system satisfies a property for all choices of lower and upper time bounds iff the underlying untimed transition system satisfies the same property under finitary weak fairness. This correspondence theorem not only establishes the adequacy of finitary fairness for distributed systems, but in addition provides a method for proving properties of timed systems whose timing is not known a priori. To summarize, finitary fairness abstracts the details of fair finite-state schedulers and the details of the independent speeds (timings) of processors with bounded drift. The parametric definition of finitary fairness also lends itself to generalizations such as computable fairness : the computable version com(F ) of a fairness assumption F is the (countable) union of all recursive safety properties that are contained in F . In a multiprogramming environment, computable fairness abstracts the details of fair computable schedulers; in a distributed environment, computable fairness abstracts the independent speeds of processors whose drift is bounded by any recursive function. Benefits of finitary fairness Verification. We address the problem of verifying that a program satisfies a property under a finitary fairness assumption fin(F ). Since fin(F ) is not !-regular, it is not specifiable in temporal logic. This, however, is not an obstacle for verification. For finite-state programs, we show that a program satisfies a temporal-logic specification under fin(F ) iff it satisfies the specification under F itself. This means that for finite-state programs, the move to finitary fairness does not call for a change in the verification algorithm. For general programs, the proof rules for verifying liveness properties are simplified by the use of finitary fairness. Suppose we wish to prove that a program terminates. To prove that all computations of a program terminate, one typically identifies a ranking (variant) function from the states of the program to the natural numbers such that the rank decreases with every transition of the program. This method is not complete for proving the termination of all fair computations. First, there may not be a ranking function that decreases at every step. The standard complete verification rule, rather, relies on a ranking function that never increases and is guaranteed to decrease eventually [LPS82, Fra86]. For this purpose, one needs to identify so-called "helpful" transitions that cause the ranking function to decrease. Second, induction over the natural numbers is not complete for proving fair termination and one may have to resort to induction over ordinals higher than !. We show that proving the termination of a program under finitary weak fairness can be reduced to proving the termination of all computations of a transformed program. The transformed program uses a new integer variable, with unspecified initial value, to represent the bound on how many consecutive times an enabled transition may be postponed. Since the termination of all computations of the transformed program can be proved using a strictly decreasing ranking function on the natural numbers, reasoning with finitary fairness is conceptually simpler than reasoning with standard fairness. Distributed consensus. A central problem in fault-tolerant distributed computing is the consensus problem, which requires that the non-faulty processes of a distributed system agree on a common output value [PSL80]. Although consensus cannot be reached in the asynchronous model if one process fails [FLP85], in practice, consensus is achieved in distributed applications using constructs like timeouts. This suggests that the asynchronous model with its standard fairness assumption is not a useful abstraction for studying fault-tolerance. One proposed solution to this problem considers the unknown-delay model (also called partially synchronous model) in which there is fixed upper bound on the relative speeds of different components, but this bound is not known a priori [DLS88, AAT97, RW92]. The asynchronous model with the finitary fairness assumption is an abstract formulation of the unknown-delay model. In particular, we prove that the asynchronous model with the finitary fairness assumption admits a wait-free solution for consensus that tolerates an arbitrary number of process failures, by showing that finitary fairness can substitute the timing assumptions of the solution of [AAT97]. Informal Motivation: Bounded Fairness Before introducing the general definition of finitary fairness (Section 3) and its applications (Sec- tions 4 and 5), we begin by motivating the finitary version of weak fairness through the intuitive concept of bounded fairness. Consider the following simple program P 0 with a boolean variable x and an integer variable y: initially repeat x := x forever k repeat y The program P 0 consists of two processes, each with one transition. The transition l complements the value of the boolean variable x; the transition r increments the value of the integer variable y. A computation of P 0 is an infinite sequence of states, starting from the initial state and such that every state is obtained from its predecessor by applying one of the two transitions. For the purpose of this example, a schedule is an infinite word over the alphabet fl; rg. Each computation of P 0 corresponds, then, to a schedule, which specifies the order of the transitions that are taken during the computation. The two processes of P 0 can be executed either by multiprogramming or in a distributed environment. Multiprogramming In a multiprogramming environment, the two processes of P 0 are scheduled on a single processor. A scheduler is a set of possible schedules. One typically requires that the scheduler is "fair"; that is, it does not shut out one of the two processes forever. Formally, a schedule is fair iff it contains infinitely many l transitions and infinitely many r transitions; a scheduler is fair iff it contains only fair schedules. be the set of fair schedules. If we restrict the set of computations of the program P 0 to those that correspond to fair schedules, then P 0 satisfies a property OE iff every computation of P 0 whose schedule is in F1 satisfies OE. For instance, under the fairness assumption F1 , the program P 0 satisfies the property that is, in any fair computation, the value of x is true in infinitely many states, and the value of y is even in infinitely many states. Note that there are computations of P 0 that correspond to unfair schedules, and do not satisfy the formula OE 1 . Thus, the fairness assumption is necessary to establish that the program P 0 satisfies the property OE 1 . The fairness requirement F1 is an abstraction of all admissible real-life schedulers, namely, those that schedule each transition "eventually." Any (non-probabilistic) real-life scheduler, however, is finite-state and therefore must put a bound on this eventuality. Consider, for instance, a round-robin scheduler that schedules the transitions l and r alternately. For round-robin schedulers, we can replace the fairness assumption F1 by the much stronger assumption F 1 that contains only two schedules, (lr) ! and (rl) ! . Under F 1 , the program P 0 satisfies the property which implies the property OE 1 . We call F 1 a 1-bounded scheduler. In general, for a positive integer k, a k-bounded scheduler never schedules one transition more than k times in a row. Formally, a schedule is k-bounded, for k - 1, iff it contains neither the subsequence l k+1 nor r k+1 ; a scheduler is k-bounded iff it contains only k-bounded schedules (similar definition is considered in [Jay88]). Let F k be the set of k-bounded schedules. The assumption F k of k-boundedness is, of course, not sufficiently abstract, because for any k, it is easy to build a fair finite-state scheduler that is not k-bounded. So let us say a schedule is bounded iff it is k-bounded for some positive integer k, and a scheduler is bounded iff it contains only bounded schedules. Clearly, every fair finite-state scheduler is bounded. In order to prove a property of the program for all implementations, then, it suffices to prove the property for all bounded schedulers. be the set of bounded schedules. If we restrict the set of computations of the program P 0 to those that correspond to bounded schedules, then P 0 satisfies a property OE iff every computation of P 0 whose schedule is in F ! satisfies OE. We call F ! the finitary restriction of the fairness assumption F1 . Three observations about F ! are immediate. First, the finitary version F ! is a proper subset of F1 ; in particular, the schedule is fair but unbounded, and therefore belongs to . Second, the set F ! itself is not a finite-state scheduler, but is the countable union of all fair finite-state schedulers. Third, F ! is again a liveness property, in the sense that a stepwise scheduler cannot paint itself into a corner [AFK88]: every finite word over fl; rg can be extended into a bounded schedule 1 . Since the finitary fairness assumption F ! is stronger than the fairness assumption F1 , a program may satisfy more properties under F ! . Consider, for example, the property where the state predicate power-of-2 (y) is true in a state iff the value of y is a power of 2. If a computation of P 0 does not satisfy OE ! , then it must be the case that the transition l is scheduled only when power-of-2 (y) holds. It follows that for every positive integer k, there is a subsequence of length greater than k that contains only r transitions. Such a schedule does not belong to F ! and, hence, the program P 0 satisfies the property OE ! under F ! . On the other hand, it is easy to construct a fair schedule that does not satisfy OE ! , which shows that P 0 does not satisfy OE ! under F1 . Multiprocessing In a distributed environment, the two processes of P 0 are executed simultaneously on two processors. While the speeds of the two processors may be different, one typically requires of a (non-faulty) processor that each transition consumes only a finite amount of time. Again, the fairness requirement F1 is an abstraction of all admissible real-life processors, namely, those that complete each transition "eventually." Again, the fairness assumption F1 is unnecessarily weak. Assume that the transition l, executed on Processor I, requires at least time ' l and at most time u l , for two unknown rational numbers ' l and u l with u l - ' l ? 0. Similarly, the transition r, It should also be noted that the set F! does not capture randomized schedulers. For, given a randomized scheduler that chooses at every step one of the two transitions with equal probability, the probability that the resulting schedule is in F! is 0. On the other hand, the probability that the resulting schedule is in F1 is 1. executed on Processor II, requires at least time ' r ? 0 and at most time u r - ' r . Irrespective of the size of the four time bounds, there is an integer k - 1 such that both k \Delta ' l ? u r and k \Delta ' r ? u l . Each computation corresponds, then, to a k-bounded schedule. It follows that finitary fairness is an adequate abstraction for speed-independent processors. It should be noted that finitary fairness is not adequate if the speeds of different processors can drift apart without bound. For this case, we later generalize the notion of finitary fairness. Finitary Fairness 3.1 Sets of infinite words An !-language over an alphabet \Sigma is a subset of the set \Sigma ! of all infinite words over \Sigma. For instance, the set of computations of a program is an !-language over the alphabet of program states. Regularity An !-language is !-regular iff it is recognized by a B-uchi automaton, which is a nondeterministic finite-state machine whose acceptance condition is modified suitably so as to accept infinite words [B-uc62]. The class of !-regular languages is very robust with many alternative characterizations (see [Tho90] for an overview of the theory of !-regular languages). In particular, the set of models of any formula of (propositional) linear temporal logic (PTL) is an !-regular language [GPSS80]. The set of computations of a finite-state program is an !-regular language. The set F1 (Section 2) of fair schedules over the alphabet fl; rg is an !-regular language (23 l - 23 r), and so is the set F k of k-bounded schedules, for every k - 1. Safety and Liveness For an !-language \Pi ' \Sigma ! , let pref (\Pi) ' \Sigma be the set of finite prefixes of words in \Pi. The !-language \Pi is a safety property (or limit-closed) iff for all infinite words w, if all finite prefixes of w are in pref (\Pi) then w 2 \Pi [ADS86]. Every safety property \Pi is fully characterized by pref (\Pi). Since a program can be executed step by step, the set of computations of a program is a safe !-language over the alphabet of program states. A safety property is !-regular iff it is recognized by a B-uchi automaton without acceptance conditions. Properties defined by temporal-logic formulas of the form 2 p, where p is a past formula of PTL, are safe and !-regular. For every k - 1, the set F k (Section 2) of k-bounded schedules is an !-regular safety property. The !-language \Pi is a liveness property iff pref that is, every finite word can be extended into a word in \Pi. The set F1 (Section 2) of fair schedules is an !-regular liveness property. Topological characterization Consider the Cantor topology on infinite words: the distance between two distinct infinite words w and w 0 is 1=2 i , where i is the largest nonnegative integer such that w The closed sets of the Cantor topology are the safety properties; the dense sets are the liveness properties. All !-regular languages lie on the first two-and-a-half levels of the Borel hierarchy: every !-regular language is in F oeffi " G ffioe . 2 There is also a temporal characterization of the first two-and-a-half levels of the Borel hierarchy [MP90]. Let p be a past formula of PTL. Then every formula of the form 2 p defines an F-set; every formula of the form 3 p, a G-set; every formula of the form 23 p, a G ffi -set; and every formula of the form 32 p, an F oe -set. For example, the set F1 of fair schedules is a G ffi -set. 3.2 The finitary restriction of an !-language Now we are ready to define the operator fin: The finitary restriction fin(\Pi) of an !-language \Pi is the (countable) union of all !-regular safety languages that are contained in \Pi. By definition, the finitary restriction of every !-language is in F oe . Also, by definition, fin(\Pi) ' \Pi. The following theorem states some properties of the operator fin. Theorem 1 Let \Pi, \Pi 0 be !-languages: 1. 2. fin is monotonic: if \Pi ae \Pi 0 then fin(\Pi) ae fin(\Pi 0 ). 3. fin distributes over intersection: fin(\Pi " Proof. The first two follow immediately from the definition of fin. Since \Pi " \Pi 0 is contained in \Pi as well as in \Pi 0 , from the monotonicity, we have To prove the inclusion From the definition of fin, there exist !-regular safety properties 1 . The class of safety properties is closed under intersection, and so is the class of !-regular languages. Hence, 1 is an !-regular safety property. Since w The following proposition formalizes the claims we made about the example in Section 2. It also shows that the finitary restriction of an !-regular language is not necessarily !-regular. Proposition 2 Let F1 be the set of fair schedules from Section 2, and let F ! be the set of bounded schedules. Then F ! is the finitary restriction of F1 (that is, neither !-regular nor safe. Proof. Recall that F is the union of !-regular safety properties contained in F1 . Each F k is an !-regular safety property and F k ae F1 . Hence, Now consider an !-regular safety property G contained in F1 . Suppose G is accepted by a B-uchi automaton MG over the alphabet fl; rg. Without loss of generality, assume that every state of MG is reachable from some initial state, and every state is an accepting state (since G is a safety property). We wish to prove that if MG has k states then G ' F k . Suppose not. Then there is a word w such that MG accepts w and w contains consecutive symbols of the same type, say l. Thus Since MG has only k states, it follows 2 The first level of the Borel hierarchy consists of the class F of closed sets and the class G of open sets; the second level, of the class G ffi of countable intersections of open sets and the class F oe of countable unions of closed sets; the third level, of the class F oeffi of countable intersections of F oe -sets and the class G ffioe of countable unions of G ffi -sets. that there is a state s of MG such that there is a path from the initial state to s labeled with w 0 l i for some 0 - i - k, and there is a cycle that contains s and all of whose edges are labeled with l. This implies that MG accepts the word w which is not a fair schedule, a contradiction to the inclusion G ' F1 . Observe that, for each k - 1, the schedule l k can be extended to a word in F k , and hence implying that F ! is not closed under limits, that is, F ! is not a safety property. Now we prove that F ! is not !-regular. Suppose F ! is !-regular. From the closure properties of !-regular languages, the set of unbounded fair schedules is also !-regular. We know that G is nonempty (it contains the schedule From the properties of !- regular languages it follows that G contains a word w such that for two finite words contains at least one l and one r symbol. This means that, for a contradiction to the assumption that w 62 F ! . In other words, although F ! is a countable union of safety properties that are definable in PTL, F ! itself is neither a safety property nor definable in PTL. To define F ! in temporal logic, one would need infinitary disjunction. In general, the operator fin does not preserve liveness also. That is, it may happen that pref However, when applied to !-regular properties, liveness is preserved. Theorem 3 If \Pi be an !-regular language then pref Proof. Since fin(\Pi) ' \Pi and pref is monotonic, pref (fin(\Pi)) ' pref (\Pi). To prove the inclusion pref (\Pi) ' pref (fin(\Pi)), suppose \Pi is !-regular language over \Sigma, and consider w 2 pref (\Pi). From !-regularity of \Pi, it follows that there is a word w 0 2 \Pi such that w 2 for finite words . The language containing the single word w 0 is !-regular, safe, and contained in \Pi. Hence, This immediately leads to the following corollary: Corollary 4 If \Pi is an !-regular liveness property then fin(\Pi) is live. Observe that the language F1 is !-regular and live, and hence, F ! is also live: pref This means that when executing a program, the fairness requirement F ! , just like the original requirement F1 , can be satisfied after any finite number of steps. The operator fin is illustrated on some typical languages below: 9k such that every subsequence of length k has some pg; 9k such that every subsequence with k p's has some qg. 3.3 Transition systems From standard fairness to finitary fairness Concurrent programs, including shared-memory and message-passing programs, can be modeled as transition systems [MP91]. A transition system P is a triple (Q; a set of states, T is a finite set of transitions, and Q 0 ' Q is a set of initial states. Each state q 2 Q is an assignment of values to all program variables; each transition - 2 T is a binary relation on the states (that is, - ' Q 2 ). For a state q and a transition - , let be the set of -successors of q. A computation q of the transition system P is an infinite sequence of states such that q 0 2 Q 0 and for all i - 0, there is a transition - 2 T with q for the set of computations of P . The set \Pi(P ) is a safe !-language over Q. If Q is finite, then \Pi(P ) is !-regular. A transition - is enabled at the i-th step of a computation q iff -(q i ) is nonempty, and - is taken at the i-th step of q iff q loss of generality, we assume that the set of program variables contains for every transition - 2 T a boolean variable enabled(-) and a boolean variable taken(- ). Let the scheduling alphabet \Sigma T be the (finite) set of interpretations of these boolean variables, that is, \Sigma T is the power set of the set fenabled(-); taken(-) j - 2 Tg. Given a computation q of P , the schedule oe(q) of q is the projection of q to the scheduling alphabet. The set of schedules of P , then, is a safety property over \Sigma T . A fairness requirement F for the transition system P is an !-language over the finite scheduling alphabet \Sigma T . The fairness requirement restricts the set of allowed computations of the program. In general, F is an !-regular liveness property [AFK88]. The requirement of liveness ensures that, when executing a program, a fairness requirement can be satisfied after any finite number of steps. In particular, the requirement of weak fairness WF for P is the set of all infinite words w such that for every transition - 2 T , there are infinitely many integers i - 0 with taken(-) 2 w i or no transition is enabled forever without being taken. It is specified by the following The requirement WF is !-regular and live. The requirement of strong fairness SF for P is the set of all infinite words w T such that for every transition - 2 T , if there are infinitely many steps there are infinitely many steps no transition is enabled infinitely often without being taken. It is a stronger requirement than the weak fairness (SF ae WF ), and is specified by the formula - The weak-fairness requirement WF is a G ffi -set; the strong-fairness requirement SF is neither in G ffi nor in F oe , but lies in F oeffi " G ffioe . Since both sets are !-regular, their finitary restrictions fin(WF) and fin(SF ), which belong to F oe , are again live (Corollary 4). The next theorem (a generalization of Proposition 2) shows that the finitary restriction of weak and strong fairness coincide with the appropriate notions of bounded fairness. Define a schedule T to be weakly-k-bounded , for a nonnegative integer k, iff for all transitions - of P , - cannot be enabled for more than k consecutive steps without being taken, that is, for all integers i - 0, there is an integer j, or enabled(-) 62 w j . A schedule is weakly-bounded if it is weakly-k-bounded for some k - 0. Similarly, a schedule w is strongly-k-bounded iff for all transitions - , if a subsequence of w contains k distinct positions where - is enabled, then it contains a position where - is taken, that is, for all integers for A schedule is strongly-bounded if it is strongly-k-bounded for some k - 0. Theorem 5 Let P be a transition system with the transition set T and the weak and strong fairness requirements WF and SF . For all infinite words w over \Sigma T , w and w 2 fin(SF) iff w is strongly-bounded. Proof. We will consider only the weak fairness. The set of weakly-k-bounded schedules, for a fixed k, is defined by the formula where wf (-) stands for the disjunction taken(-:enabled(- ). It follows that the set of weakly-k- bounded schedules is safe and !-regular, for all k - 0. Thus, every weakly-bounded schedule is in Now consider an !-regular safety property G contained in WF . Suppose G is accepted by a B-uchi automaton MG over the alphabet \Sigma T . Without loss of generality, assume that every state of MG is reachable from its initial state, and every path in MG is an accepting path. It suffices to prove that if MG has k states then every schedule in G is weakly-k-bounded. Suppose not. Let us say that a symbol of \Sigma T is weakly-unfair to a transition - , if it contains enabled(-) and does not contain taken(- ). From assumption, there is a word w accepted by MG and a transition - such that w contains k consecutive symbols all of which are weakly-unfair to - . Since MG has only states, it follows that there is a cycle in MG all of whose edges are labeled with symbols that are weakly-unfair to - . This implies that MG accepts a schedule that is not weakly-fair to - , a contradiction to the inclusion G ' WF . Observe that for the example program P 0 of Section 2, rg, the propositions enabled(l) and are true in every state, and the fairness requirement F1 equals both WF and SF . This implies that Corollary 6 If P is a transition system with at least two transitions, then both fin(WF ) and fin(SF) are neither !-regular nor safe. A computation q of the transition system P is fair with respect to the fairness requirement F We write \Pi F (P ) for the set of fair computations of P . A specification \Phi for the transition system P is a set of infinite words over the alphabet Q. The transition system P satisfies the specification \Phi under the fairness requirement F iff \Pi F (P ) ' \Phi. If we prove that P satisfies \Phi under the fairness assumption F , then P satisfies \Phi for all implementations of F ; if we prove that P satisfies \Phi under the finitary restriction fin(F ), then P satisfies \Phi for all finite-state implementations of F . In Section 4, we show that proving the latter is conceptually simpler than proving the former. 3.4 Timed transition systems From timing to finitary fairness Standard models for real-time systems place lower and upper time bounds on the duration of delays [HMP94, MMT91]. Since the exact values of the time bounds are often not known a priori, it is desirable to design programs that work for all possible choices of time bounds. It has been long realized that the timing-based model with unknown delays is different from, and often more appropriate than, the asynchronous model (with standard fairness) [DLS88, AAT97, RW92]. We show that the unknown-delay model is equivalent to the asynchronous model with finitary fairness. Real-time programs can be modeled as timed transition systems [HMP94]. A timed transition system P ';u consists of a transition system two functions ' and u from the set T of transitions to the set Q?0 of positive rational numbers. The function ' associates with each transition - a lower bound ' - ? 0; the function u associates with - an upper bound u - . The interleaving semantics of transition systems is extended to timed transition systems by labeling every state of a computation with a real-valued time stamp. A time sequence t is an infinite nondecreasing and unbounded sequence of real numbers. For t to be consistent with a given computation q of the underlying transition system P , we require that a transition - has to be enabled continuously at least for time ' - before it is taken, and it must not stay enabled continuously longer than time u - without being taken. Note that if a transition - is enabled in all states q n , for i - n - j, and - is not taken in all states q n , for then it has been continuously enabled for t Then, the time sequence t is consistent with the computation q iff for every transition - 2 T , [Lower bound ] if taken(-) 2 q j , then for all steps i with and taken(-) 62 q [Upper bound ] if enabled(-) 2 q k for all steps k with i - k - j, and taken(-) 62 q k for all steps k with A timed computation (q; t) of the timed transition system P ';u consists of a computation q of P together with a consistent time sequence t. The first component of each timed computation of P ';u is an untimed computation of P ';u . We write \Pi ';u (P ) for the set of untimed computations of P ';u . In general, \Pi ';u (P ) is a strict subset of \Pi(P ); that is, the timing information ' and u plays the same role as fairness, namely, the role of restricting the admissible interleavings of enabled transitions. If Q is finite then, like \Pi(P ), \Pi ';u (P ) is also !-regular [AD94] (but not necessarily safe). While the timed computations are required for checking if a system satisfies a specification that refers to time, the untimed computations suffice for checking if the system satisfies an untimed specification \Phi ' timed transition system P ';u satisfies the specification \Phi iff \Pi ';u (P In the unknown-delay model, we do not know the bound functions ' and u, but rather wish to prove that a transition system P satisfies the specification \Phi for all possible choices of bound functions; that is, we wish to prove that the union [ ';u \Pi ';u (P ) is contained in \Phi. The following theorem shows that in order to verify a system in the unknown-delay model, it suffices to verify the system under finitary weak fairness; that is, the union [ ';u \Pi ';u (P ) is same as the set \Pi fin(WF) (P ). Theorem 7 Let P be a transition system with the set T of transitions, let WF be the weak fairness requirement for P , and let q be a computation of P . Then q 2 \Pi fin(WF) (P ) iff for some function ' and u from T to Q?0 , q 2 \Pi ';u (P ). Proof. Consider a weakly-bounded computation q 2 \Pi fin(WF) (P ). From Theorem 5, there is a nonnegative integer k such that the schedule corresponding to q is weakly-k-bounded. Let the bound functions be defined as ' . Consider the time sequence increases by 1 at every step, it is clear that the lower bound requirement is trivially satisfied. Since q is weakly-k-bounded, no transition is enabled for more than k consecutive steps (and hence, for more than k being taken. Thus, the consistency requirements are satisfied, and (q; t) is a timed computation of P ';u , implying To prove the converse, suppose q 2 \Pi ';u (P ) for some choice of ' and u. Let t be a time sequence such that (q; t) is a timed computation of P ';u . Let ' be the (nonzero) minimum of all the lower bounds ' - , and u be the (finite) maximum of all the upper bounds u - . Let the number of transitions in T be n. Let k be an integer such that k ? nu ' . We claim that the schedule corresponding to q is weakly-k-bounded. Suppose not. Then there is a transition - and i - 0 such that - is enabled but not taken at all states q . At every step of q, taken(- 0 ) holds for some transition - 0 . Since k ? nu ' , it follows that, there is a transition - 0 such that taken(- 0 ) holds at more than u ' distinct states between q i and q i+k . Since '(- 0 , from the assumption that t satisfies the lower bound requirement, we have t implies that t violates the upper bound requirement of consistency, a contradiction. In conclusion, q is weakly-k-bounded, and hence, q 2 \Pi fin(WF) (P ). We point out that all lower bounds are, although arbitrarily small, nonzero, and all upper bounds are finite. This is necessary, and justified because we universally quantify over all choices of bound functions. We also point out that reasoning in a timing-based model with specific bound functions- i.e., reasoning about timed computations-can be significantly more complicated than untimed reasoning [HMP94]. Our analysis shows, therefore, that the verification of specifications that do not refer to time is conceptually simpler in the unknown-delay model than in the known-delay model. 3.5 The gap between finitary and standard fairness The definition of finitary fairness replaces a given !-language \Pi by the union of all !-regular safety properties contained in \Pi. While this definition seems satisfactory in practice, there are obvious mathematical generalizations. First, observe that the (uncountable) union of all safety properties contained in \Pi is \Pi itself. Not all safety properties, however, are definable by programs. We can obtain the computable restriction com(\Pi) of \Pi by taking the (countable) union of all recursive safety properties that are contained in \Pi (an !-language is recursive iff it is the language of a Turing machine). Clearly, com(\Pi) captures all possible implementations of \Pi, finite-state or not, and typically falls strictly between fin(\Pi) and \Pi. Computable fairness, however, does not have the two advantages of finitary fairness, namely, simpler verification rules and a solvable consensus problem. There are also alternatives between fin(\Pi) and com(\Pi), which capture all implementations of \Pi with limited computing power. Recall the sample program P 0 from Section 2. For every schedule in fin(\Pi), there is a bound, unknown but fixed, on how long a transition can be postponed. Suppose that we let this bound vary, and call a schedule linearly bounded iff the bound is allowed to increase linearly with time. While every bounded schedule is linearly bounded, the schedule is linearly bounded but not bounded. In general, given a function f(n) over the natural numbers, a schedule w O(f)-bounded iff there exists a constant k such that each of the two transitions l and r appears at least once in the subsequence w n w Finitary fairness, then, is O(1)-fairness. Moreover, for any fairness requirement F , we obtain a strict hierarchy of stronger fairness requirements f(F ), where f(F ) is the union of all O(f)-bounded schedulers that are contained in F . The algorithm presented in Section 5 can be modified so that it solves distributed consensus under the fairness requirement f(F ) for any fixed, computable choice of f . 4 Application: Program Verification We now consider the problem of verifying that a program satisfies a specification under a finitary fairness assumption. 4.1 Model checking If all program variables range over finite domains, then the set of program states is finite. The problem of verifying that such a finite-state program satisfies a temporal-logic specification is called model checking . Automated tools for model checking have been successfully used to check the correctness of digital hardware and communication protocols [CK96]. Here we examine the effects of finitary fairness on the algorithms that underlie these tools. Untimed systems Consider a finite-state transition system P with the state set Q. The set \Pi(P of computations of P is an !-regular safety property. Since Q is finite, we choose the scheduling alphabet to be Q itself. Let F ' Q ! be an !-regular fairness requirement, and let \Phi ' Q ! be an !-regular specification (given, say, by a PTL formula or a B-uchi automaton). The verification question, then, is a problem of language inclusion: P satisfies \Phi under F iff \Pi(P This problem can be solved algorithmically, because all involved languages are !-regular. Assuming finitary fairness, we need to check the language inclusion \Pi(P It is, however, not obvious how to check this, because fin(F ) is not necessarily !-regular (Corollary 6). The following theorem shows that finite-state verification for finitary fairness can be reduced to verification under standard fairness. Theorem 8 For all !-regular languages \Pi 1 and \Pi 2 , Proof. We have then so is fin(\Pi suppose that \Pi 1 " \Pi 2 is nonempty. Since both \Pi 1 and \Pi 2 are !-regular, so is contains a word w such that . The language containing the single word w is safe, !-regular, and contained in \Pi 1 . Hence, w 2 fin(\Pi 1 ), and also in fin(\Pi 1 implying that fin(\Pi 1 nonempty. As a corollary we obtain that for model checking under finitary fairness, we can continue to use the algorithms that have been developed to deal with standard fairness: Corollary 9 For a finite-state program P with set Q of states, an !-regular fairness requirement and an !-regular specification \Phi ' under the fairness assumption F iff satisfies \Phi under the fairness assumption fin(F ). Proof. We want to show that Let G be the !-language \Pi(P From the assumption that P is finite-state, \Pi(P ) is !-regular. Since \Phi is also !-regular, so is G. Now \Pi(P is empty (by Theorem Timed systems Consider a finite-state timed transition system P ';u with set Q of states. Suppose the specification does not refer to time at all, and is given as an !-regular specification \Phi ' Q ! . To verify that P ';u satisfies \Phi, we want to check the inclusion \Pi ';u (P This problem is solved by constructing a B-uchi automaton that recognizes the language \Pi ';u (P ) [AD94]. This method is applicable only when the bound functions ' and u are fully specified. In the parametric verification problem, the bound maps are not fully specified [AHV93]. The bounds are viewed as parameters: values of these parameters are not known, but are required to satisfy certain (linear) constraints (such as ' - 1 , ). The parametric verification problem, then, is specified by 1. a finite-state transition system 2. an !-regular specification \Phi ' 3. a set LU consisting of pairs ('; u) of functions from T to Q?0 . The verification problem is to check that, for every choice of ('; u) 2 LU , the resulting timed transition system P ';u satisfies the specification \Phi. Define (';u)2LU \Pi ';u (P Then, we want to check \Pi LU (P Theorem 8, together with Theorem 7, implies that the parametric verification problem is decidable when the set LU consists of all function pairs. If P is a finite-state transition system, \Phi is !-regular, and the parametric verification problem of checking the inclusion \Pi LU (P ) ' \Phi is decidable. The parametric verification problem, in general, is undecidable if the class LU constrains the allowed choices of the bound maps [AHV93]. For instance, if LU requires that ' then the parametric verification problem is undecidable. 4.2 Proof rules for termination We now turn to the verification of programs that are not finite-state. Since safety specifications are proved independent of any fairness assumptions, we need to be concerned only with liveness specifications. We limit ourselves to proving the termination of programs (or, equivalently, to proving specifications of the form 3 p for a state predicate p) under finitary weak fairness. It is straightforward to extend the proposed method to the verification of arbitrary temporal-logic specifications under the finitary versions of both weak and strong fairness. Total termination versus just termination system. The standard method for proving the termination of sequential deterministic programs can be adopted to prove that all computations of the (nondeter- ministic) transition system P terminate, which is called the total termination of P . Essentially, we need to identify a well-founded domain (W; OE) and a ranking (variant) function from the program states to W such that the rank decreases with every program transition. As an example, consider the rule T from [LPS82]. Rule T for proving total termination: Find a ranking function ae from Q to a well-founded domain (W; OE), and a state predicate (T2) For all states q; q Figure The rule T is complete for proving total termination; that is, all computations of a transition system P terminate iff the rule T is applicable [LPS82]. Furthermore, it is always sufficient to choose the set N of natural numbers as the well-founded domain W . Now consider the requirement that all weakly-fair computations of P terminate, which is called the just termination of P . While the rule T is obviously sound for proving just termination, it is not complete. The problem is that there may not be a ranking function that decreases with every program transition. The standard solution is to identify a ranking function that never increases, and that is guaranteed to decrease eventually. The decrease is caused by so-called "helpful" transitions, whose occurrence is ensured by the weak-fairness requirement. As an example, consider the rule J from [LPS82]. Rule J for proving just termination: Find a ranking function ae from Q to a well-founded domain (W; OE), and a set R - of state predicates, one for each transition - 2 T . Let R be the union of all R - for - 2 T . Show for all states q; q 0 2 Q and all transitions - some transition is enabled in q, then - is enabled in q. The rule J is complete for proving just termination: all weakly-fair computations of a transition system P terminate iff the rule J is applicable. Completeness, however, no longer holds if we require the well-founded domain to be the set N of natural numbers: there are transition systems for which transfinite induction over ordinals higher than ! is needed to prove just termination. An example Before we present the method for proving termination of finitary fair computations, let us consider an example. Consider the transition system P 1 of Figure 1. A state of the program P 1 is given by the values of its two variables: the location variable - ranges over f0; 1g, and the data variable x is a nonnegative integer; initially The four transitions e 1 are as shown in the figure. We want to prove that all weakly-fair computations of P 1 terminate. Initially transitions e 1 and e 2 are continuously enabled. Fairness to e 2 ensures that eventually e 4 is enabled as long as x is positive, and decrements x each time, fairness to e 4 ensures that Figure 2: Transformed program fin(P 1 ) eventually resulting in termination. To prove termination formally, we apply the rule J. As the well-founded domain, we choose the set N[f!g of natural numbers together with the ordinal !. Choose R e 1 and R e 3 to be the empty sets; a state (-; x) belongs to R e 2 belongs to R e 4 1. The ranking function is defined as: ae(0; x. The transitions e 1 and e 3 leave the rank unchanged, while e 2 and e 4 cause a decrease. The reader can check that the five premises (J1)-(J5) of the rule J are indeed satisfied. Notice that there is no bound on the number of steps before P 1 terminates. This unbounded nondeterminism is what makes the mathematical treatment of fairness difficult. Proving the termination of P 1 under finitary weak fairness-that is, proving the finitary just termination of P 1 -is conceptually simpler. Recall that for every computation in fin(WF ), there is an integer k such that a transition cannot be enabled continuously for more than k steps without being taken (Theorem 5). It follows that, under finitary weak fairness, P 1 must terminate within a bounded number of steps, where the bound depends on the unknown constant k. To capture this intuition, we transform the program P 1 by introducing the two auxiliary variables b and c. The initial value of the variable b is an unspecified nonnegative integer, and the program transitions do not change its value. The integer variable c is used to ensure that no transition is enabled for more than b steps without being taken. We thus obtain the new program fin(P 1 ) of Figure 2. The original program P 1 terminates under the finitary weak fairness assumption fin(WF) iff all computations of the transformed program fin(P 1 ) terminate. Thus, we have reduced the problem of proving the finitary just termination of P 1 to the problem of proving the total termination of That is, the simple rule T with induction over the natural numbers is sufficient to prove finitary just termination. A state of fin(P 1 ) is a tuple (-; x; c; b). To apply the rule T, we choose the set R to be the set of reachable states: (0; x; c; b) 2 R iff b. The ranking function is a mapping from R to the natural numbers defined by ae(0; c. The reader should check that every transition, applied to any state in R, causes the ranking function to decrease. Notice that the transformed program fin(P 1 ) has infinitely many initial states, but for any given initial state, it terminates within a bounded number of steps. Consequently, fin(P 1 ) does not suffer from the problems caused by unbounded nondeterminism. The tradeoff between proving just termination of P 1 , and total termination of fin(P 1 ) should be clear: while the rule J used for the former is more complex than the rule T used for the latter, the program fin(P 1 ) is more complex than Finitary transformation of a program Let us consider a general transformation for a given transition system consist of m transitions . The finitary transformation fin(P ) of the transition system P is obtained by introducing a new integer variable b, and for each transition - new integer variable c i . Thus, the state space of fin(P ) is Q \Theta N m+1 . The initial value of b is arbitrary; the initial value of each c i is 0. Thus, the set of initial states of fin(P ) is Q 0 \Theta N \Theta f0g m . For every transition - 2 T , the transition system fin(P ) contains a transition fin(-) such that 1. 2. 3. for 1 The following theorem establishes the transformation fin together with the simple rule T as a sound and complete proof method for finitary just termination. Theorem 11 A transition system P terminates under finitary weak fairness iff all computations of the transition system fin(P ) terminate. Proof. Suppose the program fin(P ) has a nonterminating computation q. Consider the projection q 0 of q on the state-space of P . From the transition rules of fin(P ), it is clear that q 0 is also a computation of P . The value of the bound variable b stays unchanged throughout q, let it be k. Furthermore, c i - k, is an invariant over the computation q for all 1 - i - m. Since for each transition - i , c i is incremented each time - i is enabled but not taken, it follows that the computation q 0 is weakly-k-bounded. Hence, P has a weakly-fair nonterminating computation. Conversely, consider a weakly-fair nonterminating computation q of P . From Theorem 5, q is weakly-k-bounded for some k. Define the sequence q 0 over the state-space of fin(P ) as follows. For all i - 0, q 0 j is the maximum nonnegative integer n such that the transition - j is enabled but not taken in all states q n 0 for It is easy to check that, since q is weakly-k-bounded, if the transition - j is enabled, but not taken, in state q i then c i Consequently, q 0 is a (nonterminating) computation of fin(P ). Thus, the language \Pi fin(WF) (P ) of finitary weak-fair computations of the transition system P is the projection of the language \Pi(fin(P )) of the transformed program fin(P ). It is known that given a transition system P and a fairness requirement F , there exists a transition system P 0 such that requires uncountably many states (the transformed program P 0 has one initial state for every fair computation in F ), and does not yield a proof principle for which well-founded induction over N is adequate. 5 Application: Distributed Consensus We consider the consensus problem in a shared-memory model where the only atomic operations allowed on a shared register are read and write. Formally, the consensus problem is defined as follows. There are n processes each with a boolean input value in i 2 f0; 1g. The process decides on the value v 2 f0; 1g by executing the statement decide(v). To model failures, we introduce a special transition fail i for each process. The transition fail i is enabled only if the Shared registers: initially: out =?, y[ 1. while out =? do 2. x[r 3. if y[r i 4. if x[r 5. else for do skip od; 7. r i := r 8. fi 9. od; 10. decide(out). Figure 3: Consensus, assuming finitary weak-fairness (program for process P i with input in i ) process P i has not yet decided on a value. When P i takes the transition fail i , all of its transitions are disabled, and P i stops participating. A solution to the consensus problem must satisfy agreement-that is, no two processes decide on conflicting values-and validity-that is, if a process decides on the value v, then v is equal to the input value of some process. Apart from these two safety requirements, we want the nonfailing processes to decide eventually: wait-freedom asserts that each process P i eventually either decides on some value or fails. Thus a process must not prevent another process from reaching a decision, and the algorithm must tolerate any number of process failures. The implicit fairness assumption in the asynchronous model is the weak-fairness requirement WF for all program transitions except the newly introduced fail i transitions. It is known that, even for there is no program that satisfies all three consensus requirements under the weak-fairness assumption WF [FLP85, LA87]. On the other hand, consensus can be solved in the unknown-delay model, where it is assumed that there is an upper bound \Delta on memory-access time, but the bound is unknown to the processes a priori and a solution is required to work for all values of \Delta [AAT97]. We show that the consensus algorithm of [AAT97] for the unknown-delay model solves, in fact, consensus under the finitary weak-fairness requirement fin(WF ). The algorithm is shown in Figure 3. The algorithm proceeds in rounds and uses the following shared data structures: an infinite two-dimensional array x[ ; 2] of bits, and an infinite array y[ ] whose elements have the value ?, 0, or 1. The decision value (i.e., the value that the processes decide on) is written to the shared bit out, which initially has the value ?. In addition, each process P i has a local register v i that contains its current preference for the decision value, and a local register r i that contains its current round number. If all processes in a round r have the same preference v, then the bit x[r; -v] is never set to 1, and consequently, processes decide on the value v in round r. Furthermore, if a process decides on a value v in round r, then y[r] is never set to the conflicting value - v, and every process that reaches round has the preference v for that round. This ensures agreement (see [AAT97] for more details of proofs). It is easy to check that if all processes have the same initial input v, then no process will ever decide on - v; implying the requirement of validity. It is possible that two processes with conflicting preferences for round r cannot resolve their conflict in round r, and proceed to round (r conflicting preferences. This happens only if both of them find y[r] =? first (line 3), and one of them proceeds and chooses its preference for the next round (line 7) before the other one finishes the assignment to y[r]. The finitary fairness requirement ensures that this behavior cannot be repeated in every round. In every finitarily fair computation, there is a bound k such that every process that has neither failed nor terminated takes a step at least once every k steps. Once the round number exceeds the (unknown) bound k, while a process is executing its for loop, all other processes are forced to take at least one step. This suffices to ensure termination. Theorem 12 The program of Figure 3 satisfies the requirements of agreement, validity, and wait- freedom under the finitary fairness assumption fin(WF ). By contrast, the program does not satisfy wait-freedom under the standard fairness assumption WF . Also observe that the algorithm uses potentially unbounded space, and therefore is not a finite-state program. The results of Section 4 imply that there is no algorithm that uses a fixed number of bounded registers and solves consensus under finitary fairness. Theorem 13 For two processes, there is no algorithm that uses finite memory, and satisfies the requirements of agreement, validity, and wait-freedom under the finitary fairness assumption fin(WF) (or equivalently, in the unknown-delay model). The unknown-delay model of [DLS88] consists of distributed processes communicating via messages, where the delivery time for each message is bounded, but is not known a priori. They establish bounds on the number of process-failures that can be tolerated by consensus protocol under various fault models. These bounds can be established using the finitary weak-fairness. Similar observation applies to the results on the session problem for the unknown-delay model [RW92]. Acknowledgments . Notions that are similar to k-bounded fairness, for a fixed k, have been defined in several places [Jay88]; the notion of bounded fairness seems to be part of the folklore, but we do not know of any published account. We thank Leslie Lamport, Amir Pnueli, Fred Schneider, Gadi Taubenfeld, and Sam Toueg for pointers to the literature and for helpful discussions. --R A theory of timed automata. Safety without stuttering. Appraising fairness in languages for distributed programming. Parametric real-time reasoning Defining liveness. On a decision method in restricted second-order arithmetic Consensus in the presence of partial synchrony. Impossibility of distributed consensus with one faulty process. On the temporal analysis of fairness. Temporal proof methodologies for timed transition systems. Communication and Synchronization in Parallel Computation. Memory requirements for agreement among unreliable asynchronous processes. Time constrained automata. A hierarchy of temporal properties. The temporal logic of reactive and concurrent systems. Reaching agreement in the presence of faults. The impact of time on the session problem. Automata on infinite objects. Verification of concurrent programs: the automata-theoretic framework --TR Fairness Safety without stuttering Consensus in the presence of partial synchrony A hierarchy of temporal properties (invited paper, 1989) Automata on infinite objects The temporal logic of reactive and concurrent systems The impact of time on the session problem Parametric real-time reasoning A theory of timed automata Temporal proof methodologies for timed transition systems Impossibility of distributed consensus with one faulty process Computer-aided verification Formal methods Time-Adaptive Algorithms for Synchronization Reaching Agreement in the Presence of Faults On the temporal analysis of fairness Impartiality, Justice and Fairness Time-Constrained Automata (Extended Abstract) Design and Synthesis of Synchronization Skeletons Using Branching-Time Temporal Logic Communication and synchronization in parallel computation

distributed consensus;modeling of asynchronous systems;fairness;program verification

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Space/time-efficient scheduling and execution of parallel irregular computations.

In this article we investigate the trade-off between time and space efficiency in scheduling and executing parallel irregular computations on distributed-memory machines. We employ acyclic task dependence graphs to model irregular parallelism with mixed granularity, and we use direct remote memory access to support fast communication. We propose new scheduling techniques and a run-time active memory management scheme to improve memory utilization while retaining good time efficiency, and we provide a theoretical analysis on correctness and performance. This work is implemented in the context of the RAPID system which uses an inspector/executor approach to parallelize irregular computations at run-ti me. We demostrate the effectiveness of the proposed techniques on several irregular applications such as sparse matrix code and the fast multipole method for particle simulation. Our experimental results on Cray-T3E show that problems large sizes can be solved under limited space capacity, and that the loss of execution efficiency caused by the extra memory management overhead is reasonable.

INTRODUCTION Considerable effort in parallel system research has been spent on time-efficient par- allelizations. This article investigates the trade-off between time and space efficiency in executing irregular computations when memory capacity on each processor is limited, since a time-efficient parallelization may lead to extra space requirements. The definition of "irregular" computation in the literature is actually not very clear. Normally in scientific computing, code with chaotic or adaptive computation and communication patterns is considered "irregular". Providing effective system support with low overhead for irregular problems is difficult and has been identified as one of the key issues in parallel system research [Kennedy 1996]. A number of projects have addressed software techniques for parallelizing different classes of irregular code [Wen et al. 1995; Das et al. 1994; Lain and Banerjee 1994; Gerasoulis Authors' addresses: T. Yang, Department of Computer Science, University of California, Santa Barbara, CA 93106. C. Fu, Siemens Pyramid, Mailstop SJ1-2-10, San Jose, CA 95134. This work was supported in part by NSF CCR-9409695, CDA-9529418, INT-9513361, and CCR-9702640, and by a DARPA subcontract through UMD (No. Z883603). A preliminary version of this article appeared in the 6th ACM Symposium on Principles & Practice of Parallel Programming (PPoPP'97). Yang and C. Fu et al. 1995; Fink et al. 1996]. This article addresses scheduling issues for parallelism that can be modeled as directed acyclic dependence graphs (DAGs) [Sarkar 1989] with mixed granularity. This model has been found useful in performance prediction and code optimization for parallel applications which have static or slowly changing dependence patterns, such as sparse matrix problems and particle simulations [Chong et al. 1995; Fu and Yang 1996b; Gerasoulis et al. 1995; Jiao 1996]. In these problems, communication and computation are irregularly interleaved with varying granularity. Asynchronous scheduling and fast communication techniques are needed for exploiting data locality and balancing loads with low synchronization costs. However, these techniques may impose extra space requirements. For example, direct memory access normally has much lower software overhead than high-level message-passing primitives. Remote addresses however need to be known at the time of performing data accesses, and thus accessible remote space must be allocated in advance. A scheduling optimization technique may prefetch or presend some data objects to overlap computation and communication, but it may require extra temporary space to hold these data objects. Therefore, using advanced optimization techniques adds difficulties in designing software support to achieve high utilization of both processor and memory resources. In this article, we present two DAG scheduling algorithms that minimize space usage while retaining good parallel time efficiency. The basic idea is to use volatile objects as early as possible so that their space can be released for reuse. We have designed an active memory management scheme that incrementally allocates necessary space on each processor and efficiently executes a DAG schedule using direct remote memory access. We provide an analysis on the performance and correctness of our techniques. The proposed techniques are implemented in the RAPID programming tool [Fu and Yang 1996a] which parallelizes irregular applications at run-time. The original schedule execution scheme in RAPID does not support incremental memory allo- cation. With sufficient space, RAPID delivers good performance for several tested irregular programs such as sparse Cholesky factorization, sparse triangular solvers [Fu and Yang 1997], and Fast Multipole Method for N-body simulation [Fu 1997]. In particular we show that RAPID can be used to parallelize sparse LU (Gaussian elimination) with dynamic partial pivoting, which is an important open parallelization problem in the literature, and deliver high megaflops on the Cray-T3D/T3E [Fu and Yang 1996b]. Another usage of RAPID is performance prediction, since the static scheduler in RAPID can predict the potential speedup for a given DAG with reasonable accuracy. For example, the RAPID sparse LU code achieves 70% of the predicted speedup [Fu and Yang 1996b]. We have found that the size of problems that RAPID can solve is restricted by the available amount of memory, which motivates us to study space optimization. It should be noted that other forms of space overhead exist, such as space for the operating system kernel, hash tables for indexing irregular objects, and task graphs. This article focuses on the optimization of space usage dedicated to storing the content of data objects. The rest of the article is organized as follows. Section 2 summarizes related work. Section 3 describes our computation and memory model. Section 4 presents space-efficient scheduling algorithms. Section 5 discusses the use T. Yang and C. Fu \Delta 3 of scheduling algorithms in the RAPID and the active memory management scheme for executing schedules. Section 6 gives experimental results. 2. RELATED WORK Most of the previous research on static DAG scheduling [Sarkar 1989; Wolski and Feo 1993; Yang and Gerasoulis 1992; 1994] does not address memory issues. A scheduling algorithm for dynamic DAGs proposed Blelloch et al. [1995] requires space usage on each processor, where S 1 is the sequential space re- quirement, p is the total number of processors, and D is the depth of a DAG. This work provides a solid theoretical ground for space-efficient scheduling. Their space model is different from ours, since it assumes a globally shared memory pool. In our model, a space upper bound is imposed to each individual processor, which is a stronger constraint. The Cilk run-time system [Blumofe et al. 1995] addresses space efficiency issues, and its space complexity is O(S 1 ) per processor, which is good in general, but too high for solving large problems if space is limited. The memory-optimizing techniques proposed in this article has space usage close to S 1 =p per processor. Another distinct difference between our work and that of Blelloch et al. [1995] and Blumofe et al. [1995] is that for RAPID, a DAG is obtained at the run-time inspector stage before its execution, thus making our scheduling scheme static. In the other models, DAGs grow on-the-fly as computation proceeds, and dynamic scheduling is preferred. There are two reasons why we use static scheduling: (1) in practice it is difficult to minimize the run-time control overhead of dynamic scheduling when parallelizing sparse code with mixed granularity on distributed memory machines; (2) the application problems we consider all have an iterative nature, so optimized static schedules can be used for many iterations. Our work uses hardware support for directly accessing remote memory, which is available in several modern parallel architectures and workstation clusters [Ibel et al. 1996; Stricker et al. 1995; Schauser and Scheiman 1995]. Advantages of direct remote memory access have been identified in fast communication research such as active messages [von Eicken et al. 1992]. Thus, we expect that other researchers can benefit from our results when using fast communication support to design software layers. 3. THE COMPUTATION AND MEMORY MODEL Our computation model consists of a set of tasks and a set of distinct data objects. Each task reads and writes a subset of data objects. Data dependence graphs (DDG) derived from partitioned code normally have three types of dependence between tasks: true dependence, antidependence, and output dependence [Poly- chronopoulos 1988]. In a DDG, some anti- or output dependence edges may be redundant if they are subsumed by other true data dependence edges. Other anti/output dependence edges can be eliminated by program transformation. This article deals with a transformed dependence graph that contains only acyclic true dependencies. An extension to the classical task graph model is that commuting tasks can be marked in a task graph to capture parallelism arising from commutative operations. The details on this parallelism model are described in Fu and Yang [1996a;1997]. We define some terms used in our task model as follows: Yang and C. Fu -A DAG contains a set of tasks, directed dependence edges between tasks, and data objects accessed by each task. Let R(T x ) be a set of objects that task T x reads and W (T x ) be a set of objects that task T x writes. -Each data object is assigned to a unique owner processor. Each processor may allocate temporary space for objects which it does not own. If a processor P x owns m, then m is called a permanent object of P x is called a volatile object of P x . Define PERM(P x ) as the set of permanent data objects on processor P x and V OLAT ILE(P x ) as the set of volatile data objects on processor P x . A permanent object will stay allocated during execution on its owner processor. -A static schedule gives a processor assignment of all tasks and an execution order of tasks on each processor. Define TA(P x ) as the set of tasks executed on processor P x . Let term "T x executed before T y on the same processor in a schedule, and let "T x ! T y " denote that task T x is executed immediately before T y . Given a DAG G and a static schedule for G, a scheduled graph is derived by marking execution edges in G (e.g., add an edge from T x to T y in G if T x ! T y ). A schedule is legal if the corresponding scheduled graph is acyclic and T x dependence edge G. -Each task T x has a weight denoting predicted computation time (called - x ) to complete this task. Each edge from T x to T y has a weight denoting predicted communication delay (called c x;y ) when sending a corresponding message between two processors. Using weight information, a static scheduling algorithm can optimize the expected parallel time of execution.3711158T 1 (a) (b) (c) Fig. 1. (a) A DAG. (b) A schedule for the DAG on 2 processors. (c) Another schedule. T. Yang and C. Fu \Delta 5 Figure 1(a) shows a DAG with 20 tasks that access 11 data objects d Each task is represented as either T i;j that reads d i and d j , and writes d j , or T j that reads and writes d j . Table I lists read and write sets of a few tasks. Parts (b) and (c) of Figure are two schedules for the DAG in (a). Permanent objects are evenly distributed between two processors. g. g. The "owner computes" rule is used to assign tasks to processors. In this example, we assume that each task and each message cost one unit of time, messages are sent asynchronously, and processor overhead for sending and receiving messages is ignored. Table I. Read and Write Sets of a Few Tasks in Figure 1(a) Notice that our task graph model does not use the SSA form (Static Single Assignment) [Cytron et al. 1991], and notice that different tasks can modify the same data object. The main reason is that in our targeted applications, a task graph derived at run-time is normally data dependent; if the SSA form were used, there would be too many objects to manage. A space/time-efficient scheduling algorithm needs to balance two conflicting goals: shortening the length of a schedule and minimizing the space requirement. If memory is not sufficient to hold all objects during the execution stage, space recycling for volatile data will be necessary. The following two strategies can be used to release the space of a volatile data object at a certain execution point: (1) the value of this object is no longer used; (2) this object is no longer accessed. As we discuss in Section 5, the second strategy reduces the overhead of data management, and we therefore choose it in our approach. For such a strategy, the space requirement is estimated as follows. Definition 1. Given a task execution order (!) on each processor, a volatile object m on a processor is called live at some task T y if it is accessed at this task or it has been accessed before and will still be accessed in the future. Formally, if the following holds, then m is live: 9T x 9T z ((T x otherwise m is called dead. Definition 2. Let size(m) be the size of data object m. For any task Tw on processor P x (i.e., Tw 2 TA(P x )), we compute the volatile space requirement at Tw on P x as m is a live volatile object at Tw Yang and C. Fu Then the memory requirement of a schedule is 8Px f In the schedule of Figure 1(b), volatile object d 3 on processor P 1 is dead after task T 3;10 , and d 5 is dead after T 5;10 . If we assume each data object is of unit size, it is easy to calculate that MEM 9. However, for the schedule in Figure 1(c), MEM REQ is 8. This is because the lifetime of volatile objects d 7 and d 3 are disjoint on P 1 , and space sharing between them is possible. 4. SPACE- AND TIME- EFFICIENT SCHEDULING Given a DAG and p processors, our scheduling approach contains two stages: -The "owner computes" rule is used to assign tasks to processors. Tasks that modify the same data objects are mapped to the same cluster, and clusters are then evenly assigned to processors. For example, given the DAG in Figure 1(a) and two processors, the set of tasks assigned to processor P 0 is g, and the set of tasks for processor P 1 is fT g. Based on the above task assign- ment, we can determine permanent and volatile data objects on each processor -Tasks at each processor are ordered, following DAG dependence edges. We focus on the optimization of task ordering. Previously, an ordering algorithm called [Yang and Gerasoulis 1992] has been proposed. It is time efficient, but it may require extra space to hold volatile objects in order to aggressively execute time-critical tasks. In this section, we propose two new ordering algorithms called MPO and DTS. The idea is to have volatile objects referenced as early as possible once they are available in the local memory. This shortens the lifetime of volatile objects and potentially reduces the memory requirement on each processor. 4.1 The RCP Algorithm We briefly discuss this algorithm; a detailed description can be found in Yang and Gerasoulis [1992]. This heuristic orders tasks by simulating the execution of tasks. Let an exit task be a task with no children in a given DAG. The time priority of a task T x , called TP (T x ), is the length of the longest path from this task to an exit task. Namely, TP (T x ) is - x if it is an exit task. Otherwise, Ty is a child of During simulated execution a task is called ready if its parents have been selected for execution and the needed data objects can be received at this point. At each scheduling cycle, each processor selects the ready task with the highest time priority. T. Yang and C. Fu \Delta 7 (1) while there is at least one un-scheduled task (2) Find the processor Px that has the earliest idle time; (3) Schedule the ready task Tx that has the highest priority on processor Px; Update the ready task list on each processor; end-while Fig. 2. The RCP algorithm. The RCP algorithm is summarized in Figure 2. Lines (4) and (5) do not exist in this description, because we want to illustrate the difference between RCP and MPO. 4.2 MPO: Memory-Priority-Guided Ordering The MPO ordering algorithm is listed in Figure 3. The difference between MPO and RCP is (1) the priority used in MPO for selecting a ready task is the size of total object space that has been allocated divided by the size of total object space needed to execute this task (if there is a tie, the RCP time priority (TP ()) is used to break the tie); (2) MPO needs to estimate the total object space allocated at each scheduling cycle and update the memory priority of tasks (lines (4) and (5)). (1) while there is at least one un-scheduled task (2) Find the processor Px that has the earliest idle time; (3) Schedule the ready task Tx that has the highest priority on processor Px; Allocate all volatile objects that Tx uses and that have not been allocated yet on processor Px; (5) Update the priorities of Tx's children and siblings on processor Px; Update the ready task list on each processor; end-while Fig. 3. The MPO algorithm. Example. Figure 1(c) shows a schedule produced by MPO while (b) is produced by RCP. The ordering difference between Figure 1(b) and (c) starts from time 6 on processor 1. At that time, T 7;8 is selected by RCP while T 3;10 is chosen by MPO. As illustrated in Figure 4, there are three ready tasks on processor 1 at time selects T 7;8 because it has the longest path to an exit task (the path is T length 4). For MPO, T 3;10 has the highest space priority (1) because d 3 and d 10 are available locally at time 6. T 7;8 's space priority is 0.5 because the space for d 7 has not been allocated before time 6, and we assume each object has a unit size. As a result, the MPO schedule requires less memory, but longer parallel time. Algorithm Complexity. A potentially time consuming part is the update of space priorities of unscheduled tasks. At line (5), it is sufficient to update the space priorities of the children and the siblings of the newly scheduled task, because only those tasks are possible candidates for ready tasks in the next round. The space priorities of the children of these candidate tasks will be updated later after the candidate tasks are scheduled. We use e in (x) and e out (x) to denote the number Yang and C. Fu (0.5, Ready task list on Proc 1: (a) (b) Fig. 4. The scheduling scenario at time 6. Numbers in parentheses next to the ready tasks are the MPO space priority and RCP time priority at that time. (a) The remaining unscheduled tasks. (b) A partial schedule at time 6. of incoming edges and outgoing edges of task T x , respectively. The complexity for line (5) is O( is the number of tasks and e is the number of dependence edges. In term e in (x) v, we have used v to bound the number of children that any parent of T x may have. Additionally, the complexities for lines (2), (3), (4), and (6) are v log p, v log v, m, and vp log v, respectively. Here, p is the total number of processors, and m is the total number of data objects. Since p is usually small compared to v, the total time complexity of MPO is O(ve 4.3 DTS: Data-Access-Directed Time Slicing DTS is more aggressive in space optimization and is thus intended for cases when memory usage is of primary importance. Its design is based on the fact that memory usage of a processor can be improved if each volatile object has a short life span. In other words, the time period from allocation to deallocation is short. The basic idea of DTS is to slice computation based on data access patterns of tasks, so that all tasks within the same slice access a small group of volatile objects. Tasks are scheduled on physical processors slice by slice, and tasks within each slice are ordered using dependence and critical-path information. Algorithm. For a given DAG in which a set of tasks operates on a set of data objects we describe the steps of the DTS algorithm as follows. 1: Construct a data connection graph (DCG). Each node of a DCG represents a distinct data object, and each edge represents a temporal order of data access during computation. A cycle may occur if accesses of two data objects are interleaved. For simplicity, we use the same name for a data object and its corresponding data node unless it will cause confusion. To construct a DCG, the following rules are applied based on the original DAG. -If a task T x 2 V uses but does not modify data object d i , T x is associated with data object node d i . T. Yang and C. Fu \Delta 9 -If T x modifies more than one objects, it may use those objects during compu- tation, and it does not use other objects, then T x is associated with each of those modified objects. -It is possible that a task is associated with multiple data nodes. In that case, doubly directed edges are added among those data nodes to make them strongly connected. -A directed edge is added from data node d i to data node d j if there exists a task dependence edge (T x ; T y ) such that T x is associated with data node d i and T y is associated with data node d j . The last two rules reflect the temporal order of data access during computation. -Step 2: Derive strongly connected components from a DCG, the edges between the components constitute a DAG. A task only appears in one component. Each component is associated with a set of tasks that use/modify data objects in this component, and is defined as one slice. All tasks in the same slice will be considered for scheduling together. At run-time, each processor will execute tasks slice by slice, following a topological order of slices imposed by dependencies among corresponding strongly connected components. It should be noted that a topological order of slices only imposes a constraint on task ordering. A processor assignment of tasks following the "owner computes" rule must be supplied before using DTS to produce an actual schedule. -Step 3: Use a priority based precedence scheduling approach to generate a DTS schedule from the slices derived at Step 2. Priorities are assigned to tasks based on the slices they belong to. For two ready tasks in the same slice, the task with a higher critical-path priority is scheduled first. If there is a ready task that has a slice priority lower than some other unscheduled tasks on the same processor, this task will not be scheduled until all the tasks that have higher slice priorities on this processor are scheduled. Using such a priority can guarantee that tasks are executed according to the derived slice order. Example. Figure 5 shows an example of DTS ordering for the DAG in Figure 1(a). Part (a) is the DCG, and we mark each node with the corresponding data name. Tasks within a rectangle are associated with a corresponding data object. Since this DCG is acyclic, each data node is a maximal strongly connected component and is treated as one slice. A topological ordering of these nodes produces a slice Each processor will execute tasks following this slice order, as shown in part (b), where slices are marked numerically to illustrate their execution order. The memory requirement MEM REQ is 7, compared to 9 in Figure 1(b) produced by RCP and 8 in Figure 1(c) by MPO. On the other hand, the schedule length increases from RCP through MPO to DTS, because less and less critical path information is used. Algorithm Complexity. For Step 1, the complexity of deriving the access pattern of each task (i.e., read and/or write access to each data object) is O(e log v); the complexity of mapping tasks to data nodes is O(v); and the complexity of generating edges between data nodes is O(e log m), because a check is needed to prevent duplicate edges from being added. Thus, the total complexity for Step 1 is v). For Step 2, deriving strongly connected components costs Yang and C. Fu3711158t1416 d d d d d d357 slice 1 slice 2 slice 4 slice 5 slice 6 slice 1 slice 2 slice 3 slice 4 slice 5 slice 6 slice 7 (b) (a) Fig. 5. (a) A sample DCG derived from a DAG. (b) A DTS schedule for the DAG on 2 processors. generating precedence edges among slices costs O(m log m); and the topological sorting of slices costs O(m+ e). Therefore, the total complexity of Step 2 is O(m m). The cost of Step 3 is O(v log v + e). This gives an overall complexity for DTS O(e(log v is the number of tasks, e is the number of edges, and m is the number of data objects in the original DAG. Space Efficiency. DTS can lead to good memory utilization; the following theorem gives a memory bound for DTS. First a definition is introduced. Definition 3. Given any processor assignment R for tasks, the volatile space requirement of slice L on processor P x , denoted as VPx (R; L), is defined as the amount of space needed to allocate for the volatile objects used in executing tasks of L on P x . The maximum volatile space requirement for L under R is then defined as Assuming that a processor assignment R of tasks using the "owner computes" rule produces an even distribution of data space for permanent data objects among processors, we can show the following results. Theorem 1. Given a processor assignment R of tasks and a DTS schedule on processors with slices ordered as L 1 , this schedule is executable with space usage per processor, where S 1 is the sequential space complexity and PROOF: First of all, since R leads to an even distribution of permanent objects, the permanent data space needed on each processor is S 1 =p. Suppose that a task T. Yang and C. Fu \Delta 11 T x in slice L i needs to allocate space for a volatile object d. If there should be enough space for d according to the definition of h. If i ? 1, then we claim that all the space allocated to the volatile data objects associated with slices L 1 can be freed. Therefore the extra h space on each processor will be enough to execute tasks in L i . Now we need to show the above claim is correct. Suppose not, and there is a volatile data object d 0 that cannot be deallocated after slice L i\Gamma1 , and d 0 is associated with L a , a ! i. Then there is at least one task T y 2 L j , j - i, that uses d 0 . If T y modifies d 0 , then d 0 is a permanent data object. If T y does not modify d 0 , then according to the DTS algorithm, T y should belong to slice L a instead of L j . Thus there is contradiction. 2 If a DCG is acyclic, then each data node in the DCG constitutes a strongly connected component. Therefore, each slice is associated with only one data object, and this implies that the h defined in Theorem 1 will be the size of the largest data object. Thus, we have the following corollary. Corollary 1. If the DCG of a task graph is acyclic and the maximum size of an object is of unit 1, the DTS produces a schedule which can be executed on processors using S 1 per processor, where S 1 is the sequential space complexity. We can apply Theorem 1 and Corollary 1 to some important application graphs. For the 1D column-block-based sparse LU DAGs in Fu and Yang [1996b], a matrix is partitioned into a set of sparse column blocks, and each task T k;j uses one sparse column block k to modify column block j. Figure 1 is actually a sparse LU graph. DTS produces an acyclic DCG for a 1D column-block-based sparse LU task graph as illustrated in Figure 5(a). Let w be the maximum space needed for storing a column block. According to Corollary 1, each processor will at most need w volatile space to execute a DTS schedule for sparse LU. For the 2D block-based sparse Cholesky approach described in Fu and Yang [1997], a matrix A is divided into N sparse column blocks, and a column block is further divided into at most N submatrices. The submatrix with index (i; j) is marked as A i;j . A Cholesky task graph can be structured with N layers. Layer represents an elimination process that uses column block k to modify column blocks from k + 1 to N . More specifically, the Cholesky factor computed from the diagonal block A k;k will be used to scale all nonzero submatrices within the kth column block, i.e., A i;k where Those nonzero submatrices will then be used to update the rest of the matrix, i.e., A i;j where update tasks at step k belong to the same slice associated with data objects A i;k where . Hence, the extra space needed to execute a slice is the summation of the submatrices in column block k. According to Theorem 1, a DTS schedule for the Cholesky can execute with S 1 =p +w space on each processor where w is again the maximum space needed to store a column block. The above results are summarized in the following corollary. Normally, w - and a DTS schedule is space efficient for these two problems. Corollary 2. For 1D column-block-based sparse LU task graphs and 2D block-based sparse Cholesky graphs, a DTS schedule is executable using S 1 Yang and C. Fu space 1 be for i=2 to k do Merge L i to L 0 space else space Fig. 6. The DTS slice-merging algorithm. per processor, where w is the size of the largest column block in the partitioned input matrix. Further Optimization. If the available memory space for each processor is known, say AV AIL MEM , the time efficiency of the DTS algorithm can be further optimized by merging several consecutive slices if memory is sufficient for those slices, and then applying the priority-based scheduling algorithm on the merged slices. Assuming there are k slices and a valid slice order is for a given task assignment R, the merging strategy is summarized in Figure 6. A set of new slices will be generated. Since calculating memory requirements for all slices takes O(e log m) time, the complexity of the merging process is O(v+e log m). It can be shown that the merging algorithm above produces an optimal solution for a given slice ordering. Theorem 2. Given an ordered slice sequence the slice-merging algorithm in Figure 6 produces a solution with the minimum number of slices. PROOF: The theorem can be proven by contradiction. Let the new slice sequence produced by the algorithm in Figure 6 be be an optimal sequence where t ? u. Each E or F slice contains a set of consecutive L slices from L 1 . The merging algorithm in Figure 6 groups as many of the first L slices as possible, H(R;E 1 We can take some of original L slices from slice F 2 , add them to F 1 so that the new F 1 is identical to E 1 . Let new F 2 be called 2 . Thus we can produce an optimal sequence We can apply the same transformation to F 0 2 by comparing it with Finally we can transform sequence F 1 another optimal sequence That is a contradiction, since this new sequence cannot completely cover all L slices unless slices are empty. 2 It should be noted that the optimality of the above merging algorithm is restricted for a given slice ordering. An interesting topic is to study if there exists a slice- sequencing algorithm which follows the partial order implied by the given DCG and leads to the minimum number of slices or the minimum parallel time. It can be shown [Tang 1998] that a heuristic using bin-packing techniques can be developed, and the number of slices is within a factor of two of the optimum. T. Yang and C. Fu \Delta 13 5. MEMORY MANAGEMENT FOR SCHEDULE EXECUTION In this section, we first briefly describe the RAPID run-time system to which our scheduling techniques are applied. Then, we discuss the necessary run-time support for efficiently executing DAG schedules derived from the proposed scheduling algorithms. 5.1 The RAPID System Dependence transformation analysis Dependence Task scheduling clustering & User specification: tasks, data objects, data access patterns. Data dependence graph (DDG) complete task graph Iterative asynchronous Task assignments, data object owners schedules and execution Fig. 7. The process of run-time parallelization in RAPID. Figure 7 shows the run-time parallelization process in RAPID. Each circle is an action performed by the system, and boxes on either side of a circle represent the input and output of the action. The API of RAPID includes a set of library functions for specifying irregular data objects and tasks that access these objects. At the inspector stage depicted in the left three circles of Figure 7, RAPID extracts a DAG from data access patterns and produces an execution schedule. At the executor stage (the rightmost circle), the schedule of computation is executed it- eratively, since targeted applications have such an iterative nature. For example, sparse matrix factorization is used extensively for solving a set of nonlinear differential equations. A numerical method such as Newton-Raphson iterates over the same sparse dependence graph derived from a Jacobian matrix, and the sparse pattern of a Jacobian matrix remains the same from one iteration to another. In the nine applications studied in Karmarkar [1991], the typical number of iterations for executing the same computation graphs ranges from 708 to 16,069, and the average is 5973. Thus the optimization cost spent for the inspector stage pays off for long simulation problems. It is shown in Fu [1997] that the RAPID inspector stage for the tested sparse matrix factorization, triangular solvers, and fast multipole method (FMM) with relatively large problem sizes takes 1-2% of the total time when the schedules are reused for 100 iterations. The inspector idea can be found in the previous scientific computing research [George and Liu 1981], where inspector optimization is called "preprocessing." Compared with the previous in- spector/executor systems for irregular computations [Das et al. 1994], the executor phase in RAPID deals with more complicated dependence structures. At the executor stage, RAPID uses direct Remote Memory Access (RMA) to execute a schedule derived at the inspector stage. RMA is available in modern multiprocessor architectures such as Cray-T3E (SHMEM), Meiko CS-2 (DMA), and SCI clusters (memory mapping). With RMA, a processor can write to the memory of any other processor, given a remote address. RMA allows data transfer directly from source to destination location without buffering, and it imposes much lower overhead than a higher-level communication layer such as MPI. The use of RMA complicates the design of our run-time execution control for data consistency. Yang and C. Fu However, we find that a DAG generated in RAPID satisfies the following properties, which simplifies our design. (Distinct data objects): A task T x does not receive data objects with the same identification from different parents. (Read/write there is a dependence path either from T x to T y or from T y to T x . there is a dependence path either from T x to T y or from T y to T x . -D4 (DAG sequentialization): Tasks can be executed consistently during sequential execution following a topological sort of this DAG. Namely, if m 2 R(T x ), the value of m that T x reads from memory during execution is produced by one of T x 's parents. In general, a DAG that satisfies D1, D2, and D3 may not always be sequentializable [Yang 1993]. A DAG with the above properties is called dependence-complete. For example, DAGs discussed in Figure 1 and in Section 6 are dependence-complete. A DDG derived from sequential code can be transformed into a dependence-complete DAG [Fu and Yang 1997]. 5.2 The Execution Scheme with Active Memory Management Maintaining and reusing data space during execution is not a new research topic, but it is complicated by using low-overhead RMA-based communication, since remote data addresses must be known in advance. We discuss two issues related to executing a DAG schedule with RMA: (1) Address consistency: An address for a data object at a processor can become stale if the value for this data object is no longer used and the space for this object is released. Using a classical cache coherence protocol to maintain address consistency can introduce a substantial amount of overhead. We have taken a simple approach in which a volatile object is considered dead only if the object with the same name will not be accessed any more on that processor. In this way, a volatile object with the same name will only be allocated once at each processor. This strategy can lead to a slightly larger memory requirement, but it reduces the complexity of maintaining address consistency. The memory requirement estimated in Section 3 follows this design strategy. (2) Address buffering: We also use RMA to transfer addresses. Since address packages are sent infrequently, we do not use address buffering, so a processor cannot send new address information unless the destination processor has read the previous address package. This reduces management overhead. The execution model using the active memory management scheme is presented below. A MAP (memory allocation point) is inserted dynamically between two consecutive tasks executed on a processor. The first MAP is always at the beginning of execution on each processor. Each MAP does the following: -Deallocate space for dead volatile objects. The dead information can be statically calculated by performing a data flow analysis on a given DAG, with a complexity proportional to the size of the graph. T. Yang and C. Fu \Delta 15 MAP allocate d8 MAP allocate d1, d3 and d5 addrs. for d1 d3 and d5 d8 addr. of d3 is suspended send d3, d5 suspended d7 is MAP allocate d7 addr. of d7 send d7 send d8 (a) MAP stop no all data objects ready MAP? yes REC END RA and CQ next task send out addrs. complete computation (b) Fig. 8. (a) MAPs in executing the schedule of Figure 2(c). (b)The control flow on each processor. -Allocate volatile space for tasks to be executed after the current point in the execution chain. Assuming that are the remaining tasks on this processor, the allocation will stop after T k if there is not enough space for executing T k+1 . The next MAP will be right before T k+1 . -Assemble address packages for other processors. Address packages may differ depending on what objects are to be accessed at other processors. Figure 8(a) illustrates MAPs and address notification when executing the schedule in Figure 1(c). If the available amount of memory is 8 for each processor, then there are 2 units of memory for volatile objects on P 1 . In addition to the MAPs at the beginning of each task chain, there is another MAP right after task T 5;10 on 1 , at which space for d 3 and d 5 will be freed and space for d 7 will be allocated. The address for d 7 on P 1 is then sent to P 0 . P 0 will send the content of d 7 to P 1 after it receives the address of d 7 . Figure 8(b) shows the control flow in our execution scheme. The system has five different states of execution: (1) REC. Waiting to receive desired data objects. If a processor is in the REC state, it cannot proceed until all the objects the current task needs are available locally. (2) EXE. Executing a task. (3) SND. Sending messages. If the remote address of a message is not available, this message is enqueued. MAP. A processor could be blocked in the MAP state when it attempts to send out address packages to other processors but a previous address package has not been consumed by a destination processor. END. At this state, the processor has executed all tasks, but it still needs to clear the send queue, which might be blocked if addresses for suspended Yang and C. Fu messages are still unavailable. For the three blocking states (i.e., the states with self-cycles in Figure 8(b)), the following two operations must be conducted frequently in order to avoid deadlock and make the execution evolve quickly: RA (Read any new address package ) and (Deliver suspended messages when addresses are available). 5.3 An Analysis on Deadlock and Consistency Theorem 3. Given a DAG G and a legal schedule, execution with the active memory management is deadlock free. Namely, the system eventually executes all tasks. PROOF: We assume that communication between processors is reliable, and we prove this theorem by induction. We will use the following two facts in our proof. -Fact 1. If a deadlock situation happens, there are a few processors blocked in a waiting cycle (e.g., a circular chain) in state REC, MAP, or END. Eventually, all processors in this circular chain only do two things (RA and CQ). The space allocation and releasing activities should complete if there is any. -Fact 2. If a processor is waiting to receive a data object, the local address for this data object must have already been notified to other processors. This is because each processor always allocates space and sends out addresses of objects before using those objects. Let G 0 be the scheduled graph of G; G 0 is acyclic because the given schedule is legal. Without loss of generality, we assume that a topological sort of G 0 produces a linear task order our induction follows this order. Induction Base. T 1 must be an entry task in G 0 , i.e., a task without parents. Before the execution of T 1 on some processor called a MAP is executed. After completing space allocation, P x only has to send out the newly created addresses. If a deadlock occurs, there are a few processors involved in a circular waiting chain. When P x is blocked in state MAP, awaiting the availability of address buffers of some destination processors, it will just do RA and CQ. Since the destination processors must be in the circular chain and they are also doing RA and CQ (according to Fact 1), their address buffers should eventually be free. Then the newly created on P x can be sent out, and P x should be able to leave the MAP state. does not have any parent, all data that T 1 needs is available locally. Hence, T 1 can complete successfully. Induction Assumption. Assume that all tasks T x for 1 - x execution. Then if T k has parents in G 0 , all of them have completed execution. We show that T k will be executed successfully. Suppose not, i.e., a deadlock occurs. Let P x be T k 's processor. The state of P x can be either MAP or REC; it cannot be state END. We discuss the following two cases. Case 1. If P x is in state REC, under the induction assumption, the only reason that P x cannot receive a data object for T k is that this object has not been sent T. Yang and C. Fu \Delta 17 out from a remote processor P y . Since all T k 's parents are finished, the only cause for P y not to send a data object out is the unavailability of its remote address on P x . According to Fact 2, the address must be already sent out to P y if P x is waiting to receive the object. Hence P y will eventually read that address through operation RA (Fact 1) and deliver the message to P x . Therefore, P x can execute T k . Case 2. If P x is in state MAP, the situation is the same as the one discussed for the induction base. This processor should be able to leave the MAP state. 2 Theorem 4. Given a dependence-complete DAG G and a legal schedule, execution with the active memory management is consistent. Namely, each task reads data objects produced by its parents specified by G. PROOF: By Theorem 3, all tasks are executed by our run-time scheme. For each dependence edge (T x ; T y ), we check if T y indeed reads object m produced by T x during execution. As illustrated in Figure 9, there are two cases in which T y could read an inconsistent copy of m, and we prove by contradiction that both of them are impossible. We assume that T y is scheduled on processor P y and T x is scheduled on P x . Time Time Dependence path/edge Write/send an object Case 1 Case 2 or or Fig. 9. An illustration of the two cases in proving Theorem 4. -Case 1 (Sender-side inconsistency): If P y 6= P x , after execution of T x , P x tries to send m to P y . This message may be suspended because the destination address may not be available. Since buffering is not used, content of m on P x may be modified before it is actually sent out to P y at time t. Assume that this case is true. Let T u be the task that overwrites m on processor sent out to P y . Then T u must intend to produce m for another task T v on P x . Since execution of T u and T x happens before T v and T y , according to Property D2, there must exist a dependence path from T x to T v and from T u to T y . According to Property D3, there must exist a dependence path between T u and T x . Yang and C. Fu If there exists a dependence path from T x to T u , the order among T x ; T u , and T y during sequential execution must be T x would not be able to read the copy of m produced by T x , which contradicts Property D4. If there exists a dependence path from T u to T x , similarly we can show that T v would not be able to read m produced by T u during sequential execution, which contradicts Property D4. -Case 2 (Receiver-site inconsistency):. After object m produced by T x is successfully delivered to the local memory of P y , the content of m on P y may be overwritten by another task (called T u it at time t. Assume that this case is true. Let T v be the task assigned to P y , and this task is supposed to read m produced by T u . According to Property D1, T v 6= T y . As illustrated in Figure 9, according to Properties D2 and D3, the dependence structure among T is the same as Case 1. Similarly, we can find a contradiction. 2 6. EXPERIMENTAL RESULTS We have implemented the proposed scheduling heuristics and active memory management scheme in RAPID on Cray-T3D/T3E and Meiko CS-2. In this section we report the performance of our approach on T3E for three irregular programs: -Sparse Cholesky factorization with 2D block data mapping [Rothberg 1992; Rothberg and Schreiber 1994; Fu and Yang 1997]. The task graph has a static dependence structure as long as the nonzero pattern of a sparse matrix is given at the run-time preprocessing stage. -Sparse Gaussian Elimination (LU factorization) with partial pivoting. This problem has unpredictable dependence and storage structures due to dynamic piv- oting. Its parallelization on shared-memory platforms is addressed in Li [1996]. However, its efficient parallelization on distributed-memory machines still remains an open problem in the scientific computing literature. We have used a static factorization approach to estimate the worst-case dependence structure and storage need. In Fu and Yang [1996b], we show that this approach does not overestimate the space too much for most of the tested matrices, and the RAPID code can deliver breakthrough performance on T3D/T3E. 1 -Fast Multipole Method (FMM) for simulating the movement of nonuniformly distributed particles. Given an irregular particle distribution, the spatial domain is divided into boxes in different levels which can be represented as a hierarchical tree. Each leaf contains a number of particles. At each iteration of a particle simulation, the FMM computation consists of upward and downward passes in this tree. At the end of an iteration, a particle may move from one leaf to another, and the computation and communication weights of the DAG which represents the FMM computation may change slightly. Since the particle movement is normally slow, the DAG representing the FMM computation can be reused for many iterations. It has been found [Jiao 1996] that static scheduling can be reused for approximately 100 iterations without too much performance degradation. A We have recently further optimized the code by using a special scheduling mechanism and eliminating RAPID control overhead, and set a new performance record [Shen et al. 1998]. T. Yang and C. Fu \Delta 19 detailed description of FMM parallelization using RAPID can be found in Fu [1997]. We first examine how the memory-managing scheme impacts parallel performance when space is limited. We then study the effectiveness of scheduling heuristics in reducing memory requirements. The reason for this presentation order is that the proposed scheduling algorithms will not be effective without proper run-time memory management. This presentation order also allows us to separate the impact of run-time memory management and new scheduling algorithms. The T3E machine we use has 128MB memory per node, and the BLAS-3 GEMM routine [Dongarra et al. 1988] can achieve 388 megaflops. The RMA primitive SHMEM PUT can achieve 0.5-2-s overhead with 500MB/sec. peak bandwidth. The test matrices used in this article are Harwell-Boeing matrix BCSSTK29 arising from structural engineering analysis for sparse Cholesky, and the "goodwin" matrix from a fluid mechanics problem for sparse LU. These matrices are of medium size 2 and solvable with any one of the three scheduling heuristics so that we can compare their performance. For FMM, we have used a distribution of 64K particles. Experiments with other test cases reach similar conclusions. In reporting parallel time under different memory constraints, we manually control the available memory space on each processor to be 75%, 50%, 40%, or 25% of TOT , where TOT is the total memory space needed for a given task schedule without any space recycling. To obtain TOT , we calculate the sum of the space for permanent and volatile objects accessed on each processor and let TOT be the maximum value among all processors. 6.1 RAPID with and without Active Memory Management 7051525Performance for Sparse Cholesky Factorization on T3E #processors 70103050Performance for Fast Multipole Method on T3E #processors (a) (b) Fig. 10. Speedups without memory optimization. (a) Sparse Cholesky. (b) FMM. 2 BCSSTK29 is of dimension 13,992 and has 1.8 million nonzeros including fill-ins; goodwin is of dimension 7320 and has 3.5 million nonzeros including fill-ins. Yang and C. Fu Table II. Absolute Performance (megaflops) for Sparse LU with Partial Pivoting Matrix P=2 P=4 P=8 P=16 P=32 P=64 goodwin 73.6 135.7 238.0 373.7 522.6 655.8 RAPID without Memory Management. Figure 10 and Table II show the overall performance of RAPID on T3E without using any memory optimization for the three test programs. This version of RAPID does not recycle space at the executor stage. The results serve as a comparison basis when assessing the performance of our memory management scheme. Note that the speedups for Cholesky and FMM are compared with high-quality sequential code, and the results are consistent with the previous work [Rothberg 1992; Jiao 1996]. The speedup for Cholesky is reasonable, since we deal with sparse matrices. The speedup for FMM is high, because leaf nodes of an FMM hierarchical tree are normally computation-intensive and have sufficient parallelism. For sparse LU, since our approach uses a static symbolic factorization which overestimates computation, we only list the megaflops performance. In calculating megaflops, we use more accurate operation counts from SuperLU [Li 1996] and divide them by corresponding numerical factorization time. RAPID with Active Memory Management. Table III examines performance degradation after using active memory management. RCP is still used for task ordering, and we show later on how much improvement on space efficiency can be obtained by using DTS and MPO. The results in this table are for sparse Cholesky (BCSSTK29), sparse LU (goodwin), and FMM under different space constraints. Column "PT inc." is the ratio of parallel time increase after using our memory management scheme. The comparison base is the parallel time of a RCP schedule with 100% memory available and without any memory management overhead. Entries marked with "1" imply that the corresponding schedule is nonexecutable under that memory constraint. The results basically show the trend that performance degradation increases as the number of processors increases and the available memory space decreases, because more overhead is contributed by address notification and space recycling. However, degradation is reasonable considering the amount of memory saved. For example, the memory management scheme can save 60% of space for Cholesky, while the parallel time is degraded by 64-93%. Observe that a schedule is more likely to be executable under reduced memory capacity when the number of processors increases. This is because more processors lead to more volatile objects on each processor, which gives the memory management scheme more flexibility to allocate and deallocate space. That is why even with 40% of the maximum memory requirement, schedules with active memory management are still executable on 16 and 32 processors, while RAPID without such support fails to execute. In Table III, we also list the average number of MAPs required one each processor. The more processors are used, the fewer MAPs are required, since less space is needed to store permanent objects on each processor. Note that for FMM, execution time with active memory management is sometimes even shorter than without memory management. An explanation for this T. Yang and C. Fu \Delta 21 Table III. Performance Degradation after Using Active Memory Management #MAP PT inc. #MAP PT inc. #MAP PT inc. P=8 2.00 38.1% 3.00 42.1% 5.00 64.1% P=32 2.00 49.2% 2.94 72.7% 3.22 94.3% LU 75% 50% 40% #MAP PT inc. #MAP PT inc. #MAP PT inc. FMM 75% 50% 40% #MAP PT inc. #MAP PT inc. #MAP PT inc. P=32 2.00 -11.5% 3.00 11.5% 5.00 18.5% is that although computation associated with leaf nodes of a particle partitioning tree is intensive, it does not mix much with intensive communication incurred in the downward and upward passes. Compared with Cholesky and LU, there are more interprocessor messages in FMM during the downward and upward passes, and the insertion of memory-managing activities enlarges gaps between consecutive communication messages, which leads to less network contention. Overhead of MAP. There are three types of memory management activities that result in time increase: RA, CQ, and MAP. Through the experiments we have found that the delivery of address packages by an MAP has never been hindered by waiting for the previous content of address buffers to be consumed. Table IV reports the overhead imposed by MAPs. It is clear that the overhead is insignificant compared with the total time increase studied in Table III. However, this activity and frequent address checking/delivering operations prolong message sending and cause the execution delay of tasks on critical paths. 6.2 Effectiveness and Comparisons of Memory-Scheduling Heuristics In this subsection, we compare the memory and time efficiency of RCP, MPO, and DTS. Memory Scalability. First we examine how much memory can be saved by using MPO and DTS. We define memory scalability (or memory reduction ratio) as , where S 1 is the sequential space requirement, and S A p is the space requirement per processor for a schedule produced by algorithm A on p processors. 22 \Delta T. Yang and C. Fu Table IV. MAP Overhead in Terms of Percentage of the Total Execution Time 75% 50% 40% 75% 50% 40% 75% 50% 40% P=32 15.4% 12.0% 10.0% 10.2% 6.8% 5.0% 4.3% 3.2% 3.3% Comparison of memory requirements for sparse Cholesky #processors Memory requirement reduction ratio x: MPO Comparison of memory requirements for sparse LU #processors Memory requirement reduction ratio x: MPO (a) (b) Comparison of memory requirements for FMM #processors Memory requirement reduction ratio x: MPO (c) Fig. 11. Memory scalability comparison of the three scheduling heuristics. (a) Sparse Cholesky. (b) Sparse LU. (c) FMM. Figure 11 shows the memory reduction ratios of the three scheduling algorithms for Cholesky, LU, and FMM. The uppermost curve in each graph is for S 1 =p, which is the perfect memory scalability. The figure shows that both MPO and DTS significantly reduce the memory requirement while DTS has a memory requirement close to the optimum in the Cholesky and LU cases. This is consistent with Corollaries 1 and 2. On the other hand, RCP is very time efficient, but not memory scalable, particularly for sparse LU. For FMM, we find that the DTS algorithm results in a single slice, i.e., all tasks belong to the same slice. The reason is that there are a lot T. Yang and C. Fu \Delta 23 of dependencies among tasks, so DTS is actually reduced to RCP. Thus, this experiment shows that if we allow the complexity to increase from O(e(log(vm))+v log v) to O(ev +m), MPO can be applied to scheduling tasks within each slice instead of RCP, which further improves space efficiency. Time Difference between RCP, MPO, and DTS. We have also compared the parallel time difference among three heuristics in Tables V and VI under different memory constraints. In these two tables, if algorithm A is compared with B (i.e., A vs. B), an entry marked by "*" indicates that the corresponding B schedule is executable under that memory constraint while the A schedule is not. Mark "-" indicates that both A and B schedules are nonexecutable. Table V. Increase of Parallel Time from RCP to MPO (RCP vs. MPO). The ratio is P=4 9.6% 11.0% 11.1% * - LU 100% 75% 50% 40% 25% FMM 100% 75% 50% 40% 25% Table shows actual parallel time increase when switching from RCP to MPO. The average increase is reasonable. Sometimes MPO schedules outperform RCP schedules even though the predicted parallel time of RCP is shorter. This is because although MPO does not use as much critical-path information as RCP does, it reduces the number of MAPs needed, and this can improve execution efficiency. Furthermore, reusing an object as soon as possible potentially improves caching performance. These factors are mixed together, making actual execution time of MPO schedules competitive to RCP. DTS is aggressive in memory saving, but it does not utilize the critical-path information in computation slicing. Table VI shows time slowdown using DTS instead of MPO. It is clear that MPO substantially outperforms DTS in terms of execution time, even though DTS is more efficient in memory usage. The difference is especially significant for a large number of processors. This is because MPO optimizes both memory usage and parallel time. However, there are times when we need DTS. For instance, in the LU case with 25% available memory, the DTS schedule Yang and C. Fu Table VI. Increase of Parallel Time from MPO to DTS (MPO vs. DTS). The ratio is LU 100% 75% 50% 40% 25% P=8 43.5% 37.4% 32.2% 5.5% - is executable on 16 processors, while the MPO schedule is too space costly to run. Note that DTS space efficiency can be further improved by using MPO to schedule each slice. Slice Merging in DTS. If the available amount of memory space is known, DTS schedules can be further optimized by the slice-merging process (called DTSM ) discussed in Section 4.3. We list the time reduction ratio by using slice merging in Table VII (DTS vs. DTSM ), and the results are very encouraging. For most cases, substantial improvement is obtained. As a result, parallel time of DTS schedules with slice merging can get very close to RCP schedules. This is because merged slices give the scheduler more flexibility in utilizing critical-path information, and DTS is also effectively improving cache performance. Thus, the DTS algorithm with slice merging is very valuable when the problem size is big and the available amount of space is known. Table VII. Reduction of Parallel Time from DTS to DTSM . The ratio is P=4 6.13% 4.85% -2.90% 7.29% - LU 100% 75% 50% 40% 25% P=32 50.55% 39.96% 38.85% 34.56% 23.95% Impact on Solvable Problem Sizes. The new scheduling algorithms can help solve problems which are unsolvable with the original RAPID system which does not optimize space usage. For example, previously the biggest matrix that could be solved T. Yang and C. Fu \Delta 25 using the RAPID LU code was e40r0100, which contains 9.58 million nonzeros with fill-ins. Using the run-time active memory management and DTS scheduling algorithm, RAPID is able to solve a larger matrix called ex11 with 26.8 million nonzeros, and it achieves 978.5 megaflops on 64 T3E nodes. In terms of single-node performance, we get 38.7 megaflops per node on 16 nodes and 13.7 megaflops per node on 64 nodes. Considering that the code has been parallelized by a software tool, these numbers are very good for T3E. 7. CONCLUSIONS Optimizing memory usage is important to solve large parallel scientific problems, and software support becomes more complex when applications have irregular computation and data access patterns. The main contribution of our work is the development of scheduling optimization techniques and an efficient memory managing scheme that supports the use of fast communication primitives available on modern processor architectures. The proposed techniques integrated with the RAPID run-time system achieve good time and space efficiency. The theoretical analysis on correctness and memory performance corroborates the design of our techniques. Experiments with sparse matrix and FMM code show that the overhead introduced by memory management activities is reasonable. The MPO heuristic is competitive to the critical-path scheduling algorithm, and it delivers good memory and time efficiency. The DTS is more aggressive in memory saving; it achieves competitive time efficiency when slice merging is conducted, and its space efficiency can be further improved by incorporating MPO for slice scheduling. It should be noted that the proposed techniques are useful for semiautomatic programming tools such as RAPID. It is still challenging to develop a fully automatic system. In the future, it is interesting to study automatic generation of coarse-grained DAGs from sequential code [Cosnard and Loi 1995], extend our results for more complicated dependence structures [Chakrabarti et al. 1995; Girkar and Polychronopoulos 1992; Ramaswamy et al. 1994], and investigate use of the proposed techniques in performance engineered parallel systems [DARPA 1998]. While massively parallel distributed-memory machines will still be valuable for high-end large-scale application problems in the future (e.g., the DOE ASCI program), an extension for SMP clusters will be useful. DTS scheduling actually also improves caching performance, and the use of this result for data placement in SMPs with memory hierarchies needs further study. ACKNOWLEDGEMENTS We would like to thank Apostolos Gerasoulis, Keshav Pingali, Ed Rothberg, Vivek Sarkar, Rob Schreiber, and Kathy Yelick for their comments on this work, the anonymous referees, and Siddhartha Chatterjee, and Vegard Holmedahl for their valuable feedbacks to improve the presentation. Theorem 2 was pointed out by one of the referees. We also thank Xiangmin Jiao for his help in implementing RAPID, Jia Jiao for providing us the FMM code and test cases, and Xiaoye Li for providing the LU test matrices. --R Provably Efficient Scheduling for Cilk: An Efficient Multithreaded Runtime System. Modeling the Benefits of Mixed Data and Task Parallelism. Multiprocessor Runtime Support for Fine-Grained Irregular DAGs Automatic Task Graph Generation Techniques. Efficiently computing static single assignment form and the control dependence graph. http://www. Communication Optimizations for Irregular Scientific Computations on Distributed Memory Architectures An Extended Set of Basic Linear Algebra Subroutines. Flexible Communication Mechanismsfor Dynamic Structured Applications Scheduling and Run-time Support for Parallel Irregular Computations Sparse LU Factorization with Partial Pivoting on Distributed Memory Machines. Also as UCSB technical report TRCS97-03 Computer Solution of Large Sparse Positive Definite Systems. Scheduling of Structured and Unstructured Computation Automatic Extraction of Functinal Parallelism from Ordinary Programs. Implementing Active Messages and Split-C for SCI Clusters and Some Architectural Implications Software Support for Parallel Processing of Irregular and Dynamic Computations. A New Parallel Architecture for Sparse Matrix Computation Based on Finite Project Geometries High Performance Fortran: Problems and Progress. Sparse Gaussian Elimination on High Performance Computers. Parallel Programming and Compilers. Exploiting the Memory Hierarchy in Sequential and Parallel Sparse Cholesky Factorization Improved Load Distribution in Parallel Sparse Cholesky Factorization. Partitioning and Scheduling Parallel Programs for Execution on Multiproces- sors Experience with Active Messages on the Meiko CS-2 Elimination Forest Guided 2D Sparse LU Factorization. Decoupling Synchronization and Data Transfer in Message Passing Systems of Parallel Computers. Personal Communication. Active Messages: a Mechanism for Integrated Communication and Computation. Runtime Support for Portable Distributed Data Structures. Program Parititoning for NUMA Multiprocessor Computer Sys- tems Scheduling and Code Generation for Parallel Architectures. of Computer Science List Scheduling with and without Communication Delays. Parallel Computing DSC: Scheduling Parallel Tasks on An Unbounded Number of Processors. revised July --TR Algorithm 656: an extended set of basic linear algebra subprograms: model implementation and test programs Efficiently computing static single assignment form and the control dependence graph A new parallel architecture for sparse matrix computation based on finite projective geometries Active messages Program partitioning for NUMA multiprocessor computer systems List scheduling with and without communication delays Techniques to overlap computation and communication in irregular iterative applications Communication optimizations for irregular scientific computations on distributed memory architectures Scheduling and code generation for parallel architectures Improved load distribution in parallel sparse Cholesky factorization Provably efficient scheduling for languages with fine-grained parallelism Modeling the benefits of mixed data and task parallelism Decoupling synchronization and data transfer in message passing systems of parallel computers Run-time compilation for parallel sparse matrix computations Run-time techniques for exploiting irregular task parallelism on distributed memory architectures Elimination forest guided 2D sparse LU factorization Sparse LU factorization with partial pivoting on distributed memory machines Partitioning and Scheduling Parallel Programs for Multiprocessors Parallel Programming and Compilers Computer Solution of Large Sparse Positive Definite Automatic Extraction of Functional Parallelism from Ordinary Programs Experience with active messages on the Meiko CS-2 Flexible Communication Mechanisms for Dynamic Structured Applications Software support for parallel processing of irregular and dynamic computations Sparse gaussian elimination on high-performance computers Scheduling and run-time support for parallel irregular computations --CTR Roxane Adle , Marc Aiguier , Franck Delaplace, Toward an automatic parallelization of sparse matrix computations, Journal of Parallel and Distributed Computing, v.65 n.3, p.313-330, March 2005 Heejo Lee , Jong Kim , Sung Je Hong , Sunggu Lee, Task scheduling using a block dependency DAG for block-oriented sparse Cholesky factorization, Proceedings of the 2000 ACM symposium on Applied computing, p.641-648, March 2000, Como, Italy Heejo Lee , Jong Kim , Sung Je Hong , Sunggu Lee, Task scheduling using a block dependency DAG for block-oriented sparse Cholesky factorization, Parallel Computing, v.29 n.1, p.135-159, January

irregular parallelism;run-time support;DAG scheduling;direct remote memory access

295662

Equality-based flow analysis versus recursive types.

Equality-based control-flow analysis has been studied by Henglein, Bondorf and Jrgensen, DeFouw, Grove, and Chambers, and others. It is faster than the subset-based-0-CFA, but also more approximate. Heintze asserted in 1995 that a program can be safety checked with an equality-based control-flow analysis if and only if it can be typed with recursive types. In this article we falsify Heintze's assertion, and we present a type system equivalent to equality-based control-flow analysis. The new type system contains both recursive types and an unusual notion of subtyping. We have s t if s and t unfold to the same regular tree, and we have &bottom;t&top; where t is a function type. In particular, there is no nontrivial subtyping between function types.

Introduction Control-flow analysis is done to determine approximate sets of functions that may be called from the call sites in a program. In this paper we address an instance of the question: Question: How does flow analysis relate to type systems? Our focus is on: 1. equality-based control-flow analysis which has been studied by Henglein [9], Bondorf and J-rgensen [3], DeFouw, Grove, and Chambers [5], and others, and 2. recursive types which, for example, are present in a restricted form in Java [6], in the form of recursive interfaces where equality and subtyping is based on names rather than structure. Equality-based control-flow analysis is a simplification of subset-based control-flow analysis [16, 11, 8]. We will use the abbreviations: subset-based control-flow analysis, and equality-based control-flow analysis. 0-CFA' is also known as, simply, 0-CFA. We can illustrate the difference between 0-CFA' and 0-CFA= by considering how they analyze a call site e 1 e 2 in a functional program. Suppose -x:e is a function in that program. We want a flow analysis to express that: if -x:e becomes the result of evaluating e 1 , then flow relations are established between the actual argument e 2 and the formal argument x, and 2) between the body e and the call site e 1 e 2 . With a subset-based analysis, the flow relations are subset inclusions. This models that values flow from the actual argument to the formal argument, and from the body of the function back to the call site. With an equality-based analysis, the flow relations are equations. Thus, the flow information for the actual and formal argument are forced to be the same, and the flow information for the body and the call site are also forced to be the same. Intuitively, the equations establish a bidirectional flow of information. 0-CFA= is more approximate than 0-CFA' . Both have been implemented many times for various purposes. In general, for functional and object-oriented languages, 0-CFA' can be executed in cubic time. For programs with finite types, 0-CFA' can be executed in quadratic time [8], and specific flow-oriented questions such as "identify all functions called from only one call site" can be answered in linear time [8]. For comparison, 0- CFA= can always be executed in almost-linear time [9]. Which one of 0-CFA' and 0- CFA= is the better choice in practice? For a language like ML [10] where functions have finite polymorphic types and data may have recursive types, experiments by Heintze and McAllester [8] indicate that it is a good choice to use 0-CFA' . They implemented a variant of the quadratic-time algorithm for 0-CFA' which treated data in a much simplified way. For the problem of pointer analysis, there are algorithms which are close cousins of 0-CFA' and 0-CFA= [17]. For this problem, the condition of finite types does not hold in general. Shapiro and Horwitz [15] presented an experimental comparison of the two algorithms, and it confirms the theoretical conclusion that 0-CFA= is faster and more approximate than 0-CFA' . For an object-oriented language like Java, the condition of finite types is seldomly satisfied because of, for example, binary methods [4]. DeFouw, Grove, and Chambers [5] experimentally compared a family of flow-analysis algorithms whose time complexities are at most cubic time. Both 0-CFA= and some of its variants do well in that comparison. Ashley [2] has also presented a flow analysis with time complexity less than cubic time. It remains open how it relates to 0-CFA= . Bondorf and J-rgensen [3] implemented both 0-CFA' and 0-CFA= for Scheme as part of the partial evaluator Similix. For Scheme, the condition of finite types does not hold in general. They concluded that the two analyses have comparable precision for their application, and that 0-CFA= is much faster. In summary, 0-CFA= has in experiments proved to be a preferable alternative to 0-CFA' for many applications. Flow analyses such as 0-CFA can be formulated using constraints, see for example [11, 14]. This approach proceeds in two steps: 1) derive flow constraints from the program text, and 2) compute the least solution of the constraints. The least solution is the desired flow information. The precision of the analysis stems from the choice of constraints. For example, one choice leads to 0-CFA' , and another choice leads to 0-CFA= . The kind of flow constraints used in, for example, the paper [11] always admits a least solution. We can turn a flow analysis into a predicate which accepts and rejects programs, by extending it with safety constraints. For example, for a call site e 1 e 2 in a functional program, a safety constraint might express: "does the flow information for e 1 denote only Safety constraints do not always have a solution. They can be derived from the program text, just like flow constraints. This means that we can do a flow-based safety analysis of a program in two steps: 1) derive flow and safety constraints from the program text, and 2) decide if the constraints are satisfiable. Such a safety analysis performs a task akin to type inference, in the sense that "safe" is like "typable." Palsberg and O'Keefe [12] showed that a program can be safety checked with 0-CFA' if and only if it can be typed in Amadio and Cardelli's type system with subtyping and recursive types [1]. The proof of this connection makes explicit the close relationship between flow and subtyping. Heintze asserted in 1995 [7] that a program can be safety checked with 0-CFA= if and only if it can be typed with recursive types. This assertion is reasonable because it says that, intuitively, if we replace subset inclusions by equalities, then the need for subtyping disappears. Heintze's assertion is also consistent with the observation that both 0-CFA= and type inference with recursive types can be executed in almost-linear time. Perhaps surprisingly, Heintze's assertion is false. For example, consider the -term: The variable f is applied to both the number 0 and the function -x:x. Thus, the -term does not have a type in a type system with recursive types but no subtyping. Still, a -based safety analysis accepts this program, by assigning both f and g the empty flow set, see Section 2 for details. For another example, consider the -term: It reminds a bit of the previous example, but now f is applied to (-a:0) and (-b:-x:x). Again, the -term e 2 does not have a type in a type system with recursive types but no subtyping. For conservative flow analysis cannot assign the empty flow set to f because that flow set should at least contain (-y:0). Still, a 0-CFA= -based safety analysis accepts this program, by assigning y a flow set which contains both (-a:0) and (-b:-x:x). Given that Heintze's assertion is false, we are left with two questions: 1. which type system corresponds to 0-CFA= ?, and 2. which control-flow analysis corresponds to recursive types? Palsberg and O'Keefe's result [12] implies that E 1 and E 2 can be typed if we have both recursive types and Amadio/Cardelli subtyping. Their result also seem to indicate that adding both recursive types and all of the Amadio/Cardelli subtyping to match 0-CFA= would be overkill. Thus, to answer the first question, it makes sense to ask: how much subtyping is necessary and sufficient to match 0-CFA= ? To answer the second question we must ask: what restrictions on 0-CFA= must we impose to match recursive types? In this paper we answer the first question and we give a partial answer to the second question. We show that a program can be safety checked with 0-CFA= if and only if it can be typed with recursive types and an unusual restriction of Amadio/Cardelli subtyping. We have s - t if s and t unfold to the same regular tree, and we have is a function type. In particular, there is no non-trivial subtyping between function types. To see why non-trivial subtyping between function types is not required to match 0-CFA= , consider the program (-x:e)e 0 . Let hxi be a flow variable for the binding occurrence of x, and let [[(-x:e)e 0 ]], [[-x:e]], flow variables for the occurrences (-x:e)e 0 , -x:e, e, e 0 , respectively. If ' is a map from flow variables to flow sets, which satisfies the 0-CFA= constraints, then in particular it satisfies We can also use hxi, variables, and for a type system such as simple types where there is no non-trivial subtyping between function types, we get, among others, the following constraints on type correctness: Unification gives that a typing must satisfy the constraints: Thus, we get the same form of relationships between the types as there are between the flow sets. If we allow non-trivial subtyping between function types, then the constraints on type correctness become [12]: In particular, this opens the possibility for a non-trivial relationship: and hence These constraints are closely related to the flow constraints used in 0-CFA' [12]. We also show that if a program can be safety checked with a certain restriction of 0-CFA= , then it can be typed with recursive types. Our restriction of 0-CFA= is that all flow sets must be nonempty and consistent. Consistency means that if two functions -x:e and -y:e 0 occur in the same flow set, then the flow sets for x and y are equal, and also the flow sets for e and e 0 are equal. In slogan-form, our results read: tiny drop of subtyping. Recursive types ' The key to understanding the second result is that both empty flow sets and flow sets with two or more inconsistent functions have no counterparts in a type system with just recursive types. The restricted version of 0-CFA= does not fully match recursive types, because a program may have a type for which no flow set exists. In the next section, we present Heintze's definition of 0-CFA= , in Section 3 we present the new type system, and in Section 4 and 5 we prove our results. Our example language is a -calculus, defined by the grammar: where succ denotes the successor function on integers. Equality-Based Control-Flow Analysis Given a -term P , assume that P has been ff-converted such that all bound variables are distinct and different from the free variables. Let Var(P ) be the set of -bound variables in P . Let X P be the set of variables consisting of one variable hxi for each x 2 Var(P ). Let Y P be a set of variables disjoint from X P consisting of one variable each occurrence of a subterm e of P . (The notation ambiguous because there may be more than one occurrence of e in P . However, it will always be clear from context which occurrence is meant.) The set Abs(P ) is the set of occurrences of subterms -x:e of P . The set CL(P ) is Flow-based safety analysis of a -term P can be phrased in terms of a constraint system over the variables range over CL(P ffl For every occurrence in P of a subterm of the form 0, the constraint ffl for every occurrence in P of a subterm of the form succ e, the two constraints ffl for every occurrence in P of a subterm of the form -x:e, the constraint ffl for every occurrence in P of a subterm of the form e 1 e 2 , the constraint ffl for every occurrence in P of a -variable x, the constraint ffl for every occurrence in P of a subterm of the form -x:e, and for every occurrence in P of a subterm of the form e 1 e 2 , the constraints The last two constraints create a connection between a call site e 1 e 2 and a potential callee -x:e. Notice that two of the constraints are not equalities, but subset inclusions. This is the key reason why subtyping is needed to match this safety analysis. This constraint system mixes flow constraints and safety constraints. The safety constraints are: ffl for succ e: [[succ ffl for e 1 and the rest are flow constraints. Notice that because Int and functions cannot occur in the same flow set we have that a constraint such as has the same effect as fIntg. Denote by C(P ) the system of constraint generated from P in this fashion. Let Cmap(P ) be the set of total functions from all constraints in C(P ). We say that P is 0-CFA= safe if C(P ) is For example, consider again where we have labeled the two occurrences of f as f 1 and f 2 , for notational convenience. We have: The constraint system C(E 1 ) has the point-wise '-least solution Next, consider again: where we have labeled the occurrences of f as f 1 and f 2 , for notational convenience. The constraint system C(E 2 ) has the point-wise '-least solution etc. 3 The Type System We use v to range over type variables drawn from a countably infinite set Tv. Types are defined by the grammar: with the restriction that a type is not allowed to contain anything of the form We identify types with their infinite unfoldings under the rule: Such infinite unfolding eliminates all uses of - in types. It follows that types are a class of regular trees over the alphabet There is a subtype relation - on types: It is straightforward to show that - is a partial order. Notice that ? is a lower bound and ? is an upper bound for only the function types but not Int. A more suggestive notation might be ?! for ?, and ?! for ?. A type environment is a partial function with finite domain which maps -variables to types. We use A to range over type environments. We use the notation A[x : t] to denote an environment which maps x to t, and maps y, where y 6= x, to A(y). A type judgment has the form A ' e : t, and it means that in the type environment A, the expression e has type t. Formally, this holds when it is derivable using the rules below. Notice that there is no subsumption rule; instead subtyping can only be used in a restricted way in rules 2 and 3. We say that e is RS-typable if A ' e : t is derivable for some A; t. (RS stands for "restricted subtyping.") The type system has the subject reduction property, that is, if A ' e : t is derivable, and e beta-reduces to e 0 , then A ' e is derivable. This can be proved by straightforward induction on the structure of the derivation of A ' e : t. Here follow type derivations for the two -terms Section 1. The first type derivation uses the abbreviation: Notice the four uses of subtyping. Notice also that the only possible type for f is ?. The second derivation uses the abbreviation: A Notice that the only possible common type for both (-a:0) and (-b:-x:x) is ?. The reason why there is no subsumption rule of the form is that we want to disallow the use of subsumption immediately after a use of the rule for variables. If we add a subsumption rule, then more -terms become typable. For example, consider: If we have a subsumption rule, then we can give -y:y the type ? ! ?, we can give both -x:0 and the last occurrence of f the type ?, and it is then straightforward to complete a type derivation for E 3 . Notice that the fragment of the type derivation for the last occurrence of f is of the form: Without a subsumption rule, this type derivation is not possible. Indeed, no type derivation using rules (1)-(5) is possible. To see that, let s 1 be the type of -y:y, let s 2 be the type of f . From -y:y we have is the type of x. Moreover, from (ff) we have where u is the type of (ff ). We have hence ff). Consider now (f(-x:0)). The type of -x:0 is of the form s or ?. In both cases, it cannot be an argument of a function of type -ff:(ff ! ff). We conclude that E 3 is not RS-typable. 4 The Equivalence Result Theorem 4.1 A -term P is 0-CFA= safe if and only if P is RS-typable. We prove this theorem in two steps. Lemma 4.3 shows that if P is 0-CFA= safe, then P is RS-typable. To prove that lemma we use the technique from [13]. Lemma 4.4 shows that if P is RS-typable, then P is 0-CFA= safe. To prove that lemma we use a technique which is more direct than the one used to show a similar result, for 0-CFA' , in [12]. From Flows to Types First we consider the mapping of flows to types. Given a program P , a map ' 2 Cmap(P ), and S ' Abs(P ), we say that S is '-consistent if for all -x 1 Given a program P and ' 2 Cmap(P ), define the equation system \Gamma(P; '): ffl For each S 2 range('), let v S be a type variable, and contains the equation contains the equation there are two cases: either S is '- consistent and then \Gamma(P; ') contains the equation otherwise \Gamma(P; ') contains the equation Every equation system \Gamma(P; ') has a unique solution. To see this, notice that for every type variable, there is exactly one equation with that variable as left-hand side. Thus, intuitively, we obtain the solution by using each equation as an unfolding rule, possibly infinitely often. Lemma 4.2 If ' 2 Cmap(P is the unique solution of Proof. Support first that '(w 1 Suppose then that '(w 1 ) is '-inconsistent. From '(w 1 ) ` '(w 2 ) we then have that also Suppose finally that '(w 1 consistent. There are two cases. If '(w 2 ) is '-inconsistent, then /(v '(w 1 Lemma 4.3 If ' satisfies C(P ), is the unique solution of \Gamma(P; '), and e is a subterm of P , then we can derive A ` e : /(v '([[e]]) ). Proof. We proceed by induction on the structure of e. In the base case, consider first We have so we can derive A ` x : /(v '(hxi) ). This is the desired derivation because Consider then e j 0. We have and we can derive In the induction step, consider first We have so From the induction hypothesis we have that we can derive A ' e and we can then also derive A ' succ e Consider next e j -x:e 0 . We have f-x:e 0 g ' '([[-x:e 0 ]]), and from Lemma 4.2 we get From the induction hypothesis, we have that we can derive Thus, we can also derive A ' -x:e Finally, consider e We have '([[e 1 ]]) ` Abs(P ), and for every -x:e 0 2 '([[e 1 ]]) we have '([[e 2 From the induction hypothesis, we have that we can derive A ' There are two cases. If '([[e 1 and we can derive A ' e 1 then we use '([[e 1 ]]) ' Abs(P ) to conclude that '([[e 1 ]]) for all is '-consistent. Thus, /(v '([[e 1 ]]) we can derive A ' e 1 For example, consider again the -term: and recall the function ' 1 from Section 2 which satisfies C(E 1 ). The constraint system When we plug this into the construction in the proof of Lemma 4.3, we get the type derivation for shown in Section 3. We leave it to the reader to carry out the construction It will lead to the type derivation for shown in Section 3. From Types to Flows Next we consider the mapping of types to flows. If \Delta is the type derivation A ' define f \Delta to map types to elements of CL(P the set of occurrences -x:e of P where \Delta contains a judgment of the form A for an occurrence -x:e of P where \Delta contains a judgment of the form A for an occurrence e of P where \Delta contains a judgment of the form A 0 Lemma 4.4 If \Delta is the type derivation A ' Proof. We consider in turn each of the constraints in C(P ). For an occurrence of 0 and the constraint we have that \Delta contains a judgment of the form A 0 fIntg. For an occurrence of succ e and the constraints we have that \Delta contains judgments of the forms A fIntg. For an occurrence x and the constraint we have that \Delta contains a judgment of the form A[x For an occurrence -x:e and the constraint f-x:eg ' [[-x:e]], we have that \Delta contains judgments of the forms A There are two cases. If f-x:eg. If For an occurrence e 1 e 2 and the constraint also the constraints, for every occurrence -x:e in Abs(P ), we have that \Delta contains judgments of the forms A 0 ' where There are two cases. If and the other constraints are vacuously satisfied. If From the definition of f Concluding Remarks If we remove from Section 3 the types ?, ? and the notion of subtyping, then we get a traditional system of recursive types. Given a program P and a map ' 2 Cmap(P ), we say that ' is consistent if for all S 2 range(') we have that S is '-consistent. If we add to Section 2 the conditions: does not contain ;, and does not contain inconsistent maps, then we get a notion of flow-based safety analysis which we here will refer to as restricted- 0-CFA= safety. It is easy to modify the proof of Lemma 4.3 to show the following result. Theorem 5.1 If a -term P is restricted-0-CFA = safe, then P is typable with recursive types. Intuitively, the theorem says that if we want a flow analysis weaker than recursive types, then we can start with 0-CFA= , outlaw ;, and insist on internal consistency in all flow sets. The converse of Theorem 5.1 is false. For example, if we attempt to modify the proof of Lemma 4.4, then we run into trouble in the case e 1 e 2 , because there is no guarantee that is the type of e 1 . Such a situation arises with the program With recursive types but not subtyping, there is just one type derivation for E 4 , using the abbreviation We have It it straightforward to show that '([[x]]) 6= fIntg and '([[x]]) 6= f-x:succ(x0)g, so E 4 is therefore a counterexample to the converse of Theorem 5.1. We leave it as an open problem to find a flow analysis equivalent to recursive types. An unusual aspect of Heintze's definition of 0-CFA= is that Int and functions cannot occur in the same flow set. To allow that we might define and change the constraints from Section 2 such that the constraints for 0 and succ e become: There is a systematic way of obtaining this modified flow analysis: begin with the constraints for 0-CFA' [12] and ffl change to All other constraints remain the same. The type system that matches the modified flow analysis can be obtained by changing the type system from Section 3 such that - is the smallest reflexive and transitive relation on types where ? - t, and such that the type rules for 0 and succ e become: Notice that in this modified type system, ? is the least type and ? is the greatest type. --R Subtyping recursive types. A practical and flexible flow analysis for higher-order languages Efficient analyses for realistic off-line partial evaluation On binary methods. Fast interprocedural class analysis. The Java Language Specification. Dynamic typing. The Definition of Standard ML. Closure analysis in constraint form. A type system equivalent to flow analysis. From polyvariant flow information to intersection and union types. Fast and accurate flow-insensitive points-to analysis --TR The definition of Standard ML Control-flow analysis of higher-order languages of taming lambda Dynamic typing Subtyping recursive types Object-oriented type systems Closure analysis in constraint form A type system equivalent to flow analysis On binary methods Points-to analysis in almost linear time A practical and flexible flow analysis for higher-order languages Linear-time subtransitive control flow analysis Fast and accurate flow-insensitive points-to analysis From polyvariant flow information to intersection and union types Fast interprocedural class analysis The Java Language Specification Control-Flow Analysis and Type Systems --CTR Jens Palsberg , Mitchell Wand, CPS transformation of flow information, Journal of Functional Programming, v.13 n.5, p.905-923, September Naoki Kobayashi, Type-based useless variable elimination, ACM SIGPLAN Notices, v.34 n.11, p.84-93, Nov. 1999 Naoki Kobayashi, Type-Based Useless-Variable Elimination, Higher-Order and Symbolic Computation, v.14 n.2-3, p.221-260, September 2001 Neal Glew , Jens Palsberg, Type-safe method inlining, Science of Computer Programming, v.52 n.1-3, p.281-306, August 2004 Jens Palsberg , Christina Pavlopoulou, From Polyvariant flow information to intersection and union types, Journal of Functional Programming, v.11 n.3, p.263-317, May 2001 Michael Hind, Pointer analysis: haven't we solved this problem yet?, Proceedings of the 2001 ACM SIGPLAN-SIGSOFT workshop on Program analysis for software tools and engineering, p.54-61, June 2001, Snowbird, Utah, United States

flow analysis;type systems

295663

A new, simpler linear-time dominators algorithm.

We present a new linear-time algorithm to find the immediate dominators of all vertices in a flowgraph. Our algorithm is simpler than previous linear-time algorithms: rather than employ complicated data structures, we combine the use of microtrees and memoization with new observations on a restricted class of path compressions. We have implemented our algorithm, and we report experimental results that show that the constant factors are low. Compared to the standard, slightly superlinear algorithm of Lengauer and Tarjan, which has much less overhead, our algorithm runs 10-20% slower on real flowgraphs of reasonable size and only a few percent slower on very large flowgraphs.

INTRODUCTION We consider the problem of nding the immediate dominators of vertices in a graph. A owgraph is a directed graph r) with a distinguished start vertex , such that there is a path from r to each vertex in V . Vertex w dominates vertex v if every path from r to v includes w; w is the immediate dominator (idom) of v, denoted dominates v and (2) every other vertex x that dominates v also dominates w. Every vertex in a owgraph has a unique immediate dominator [Aho and Ullman 1972; Lorry and Medlock 1969]. Finding immediate dominators in a owgraph is an elegant problem in graph the- ory, with applications in global ow analysis and program optimization [Aho and Ullman 1972; Cytron et al. 1991; Ferrante et al. 1987; Lorry and Medlock 1969]. Lorry and Medlock [1969] introduced an O(n 4 )-time algorithm, where to nd all the immediate dominators in a owgraph. Successive improve- Some of this material was presented at the Thirtieth ACM Symposium on the Theory of Com- puting, 1998. Authors' address: AT&T Labs, Shannon Laboratory, 180 Park Ave., Florham Park, NJ 07932; Permission to make digital/hard copy of all or part of this material without fee is granted provided that the copies are not made or distributed for prot or commercial advantage, the ACM copyright/server notice, the title of the publication, and its date appear, and notice is given that copying is by permission of the Association for Computing Machinery, Inc. (ACM). To copy otherwise, to republish, to post on servers, or to redistribute to lists requires prior specic permission and/or a fee. c 2 Adam L. Buchsbaum et al. ments to this time bound were achieved [Aho and Ullman 1972; Purdom and Moore 1972; Tarjan 1974], culminating in Lengauer and Tarjan's [1979] O(m(m;n))-time algorithm; is the standard functional inverse of the Ackermann function and grows extremely slowly with m and n [Tarjan and van Leeuwen 1984]. Lengauer and Tarjan [1979] report experimental results showing that their algorithm outperforms all previous dominators algorithms for owgraph sizes that appear in practice. Reducing the asymptotic time complexity of nding dominators to O(n +m) is an interesting theoretical exercise. Furthermore, various results in compiler theory rely on the existence of a linear-time dominators algorithm; Pingali and Bilardi [1997] give an example and further references. Harel [1985] claimed a linear-time dominators algorithm, but careful examination of his abstract reveals problems with his arguments. Alstrup et al. [1997] detail some of the problems with Harel's approach and oer a linear-time algorithm that employs powerful data structures based on bit manipulation to resolve these problems. While they achieve a linear-time dominators algorithm, their reliance on sophisticated data structures adds su-cient overhead to make any implementation impractical. We present a new linear-time dominators algorithm, which is simpler than that of Alstrup et al. [1997]. Our algorithm requires no complicated data structures: we use only depth-rst search, the fast union-nd data structure [Tarjan and van Leeuwen 1984], topological sort, and memoization. We have implemented our al- gorithm, and we report experimental results, which show that, even with the extra overhead needed to achieve linear time, our constant factors are low. Ours is the rst implementation of a linear-time dominators algorithm. The rest of this article is organized as follows. Section 2 outlines Lengauer and Tarjan's approach. Section 3 gives a broad overview of our algorithm and dieren- tiates it from previous work. Section 4 presents our algorithm in detail, and Section 5 analyzes its running time. Section 6 presents our new path-compression result, on which the analysis relies. Section 7 describes our implementation, and Section 8 reports experimental results. We conclude in Section 9. 2. THE LENGAUER-TARJAN ALGORITHM Here we outline the Lengauer and Tarjan (LT) approach [Lengauer and Tarjan 1979] at a high level, to provide some details needed by our algorithm. Appel [1998] provides a thorough description of the LT algorithm. r) be an input owgraph with n vertices and m arcs. Let D be a depth-rst search (DFS) tree of G, rooted at r. We sometimes refer to a vertex x by its DFS number; in particular, x < y means that x's DFS number is less than y's. Let w that w is an ancestor (not necessarily proper) of v in D; can also denote the actual tree path. Similarly, w that w is a proper ancestor of v in D and can represent the corresponding path. For any tree (v) be the parent of v in T , and let nca T (u; v) be the nearest common ancestor of u and v in T . We will drop the subscripts and write p(v) and nca(u; v) when the context resolves any ambiguity. v) be a path in G. Lengauer and Tarjan dene P to be a semidominator path (abbreviated sdom path) if x i > v; 1 1. An sdom path from u to v thus avoids all tree vertices between u and A New, Simpler Linear-Time Dominators Algorithm 3 a,13 d,14 (a) a,13 d,14 (b) Fig. 1. (a) A owgraph G with root r. Vertex labels are augmented with their DFS numbers. (b) A DFS tree D of G. Solid arcs are tree arcs; dotted arcs are nontree arcs. Breaking D into microtrees of size no more than 3 results in four nontrivial microtrees, rooted at k, j, g, and a; the vertices of each nontrivial microtree are encircled. v. The semidominator (semi) of vertex v is there is an sdom path from u to vg: For example, consider vertex g in DFS tree D in Figure 1(a). The DFS number of g is 10. Paths (e; g), (f; g), (d; f; g), (a; d; f; g), (b; d; f; g), and (b; a; d; f; g) are all the sdom paths to g. Since b has the least DFS number over the initial vertices on these paths, To compute semidominators, Lengauer and Tarjan use an auxiliary link-eval data structure, which operates as follows. Let T be a tree with a real value associated with each vertex. We wish to maintain a forest F contained in the tree, subject to the following operations. (Initially F contains no arcs.) to F . Let r be the root of the tree containing u in F . If return r. Otherwise, return any vertex x 6= r of minimum value on the path r u. Tarjan [1979a] shows how to implement link and eval using the standard disjoint set union data structure [Tarjan and van Leeuwen 1984]. Using linking by size and path compression, n 1 links and m evals on an n-vertex tree T can be performed in O(m(m;n) + n) time. The LT algorithm traverses D in reverse DFS order, computing semidominators as follows. (Initially, semi(v) v for all For in reverse DFS order do For (w; v) 2 A do 4 Adam L. Buchsbaum et al. done link(v) done It then computes the immediate dominator for each vertex, using semidominators and the following facts, which we will also use to design our algorithm. Lemma 2.1 (LT Lem. 1). If v w then any path from v to w in G must contain a common ancestor of v and w in D. Lemma 2.2 (LT Lem. 4). For any vertex v 6= r, idom(v) semi(v). Lemma 2.3 (LT Lem. 5). Let vertices w; v satisfy w v. Then w or idom(v) idom(w). Lemma 2.4 (LT Thm. 2). Let w 6= r. Suppose every u for which semi(w) Lemma 2.5 (LT Thm. 3). Let w 6= r, and let u be a vertex for which semi(u) is minimum among vertices u satisfying semi(w) and 3. OUTLINE OF A LINEAR-TIME ALGORITHM The links and evals used by the LT algorithm make it run in O(m(m;n)) time. We can eliminate the (m; n) term by exploiting the sensitivity of to relative dierences in m and n. In particular, when m is slightly superlinear in n, e.g., m=n becomes a constant [Tarjan and van Leeuwen 1984]. 1 Our dominators algorithm proceeds roughly as follows: (1) Compute a DFS tree D of G, and partition D into regions. We discuss the partitioning in detail in Section 3.2. For now, it su-ces to consider that D is partitioned into a collection of small, vertex-disjoint regions, called microtrees. We consider separately the microtrees at the bottom of D|those that contain the leaves of D|from the microtrees comprising the interior, D 0 , of D. (2) For each vertex, determine whether its idom is in its microtree and, if so, determine the actual idom. (3) For each vertex v such that idom(v) is not in v's microtree (a) compute idom(v) by applying the LT algorithm only to vertices in D 0 or (b) nd an ancestor u of v such that and idom(u) can be computed by applying the LT algorithm only to vertices in D 0 . Partitioning D into microtrees serves two purposes. First, the subgraph induced by the microtree roots will achieve the ratio m=n necessary to reduce (m; n) to a constant. Second, when the microtrees are small enough, the number of distinct microtrees will be small compared to n +m. We can thus perform simple computations on each microtree in O(n using precomputed tables or memoization to eliminate redundant computations. log (i) n is the iterated log function: log (0) A New, Simpler Linear-Time Dominators Algorithm 5 3.1 Comparison to Previous Approaches We contrast our use of these facts to previous approaches. Harel [1985] and Alstrup et al. [1997] apply the LT algorithm to all of D, using the microtree partitioning to speed links and evals. Harel [1985] divides the entire tree D into microtrees, all of which can contain more than one vertex, and performs links and evals as described by Lengauer and Tarjan [1979] on the tree D 0 induced by the microtree roots. Alstrup et al. [1997] simplify Harel's approach [Harel 1985] by restricting nonsin- gleton microtrees to the bottom of D, leaving an upper subtree, D 0 , of singleton microtrees as we do. They then perform links and evals on D 0 using two novel data structures, as well as the Gabow-Tarjan linear-time disjoint set union result [Gabow and Tarjan 1985] and transformations to D 0 . Both algorithms use pre-computed tables to process evals on the internal microtree vertices. This approach requires information regarding which vertices outside microtrees might dominate vertices inside microtrees, to derive e-cient encodings needed by the table lookup technique. Harel [1985] presents a method to restrict the set of such outside dominator candidates. Alstrup et al. [1997] demonstrate deciencies in Harel's arguments and correct these problems, using Fredman and Willard's Q-heaps [Fredman and Willard 1994] to manage the microtrees. We apply the LT algorithm just to the upper portion, D 0 , of D. We combine our partitioning scheme with a new path compression result to show that the LT algorithm runs in linear time on D 0 . Instead of processing links and evals on internal microtree vertices, we determine, using any simple dominators algorithm, whether the dominators of such vertices are internal to their microtrees, and if so we compute them directly, using memoization to eliminate redundant computation. We process those vertices with dominators outside their microtrees without performing evals on internal microtree vertices. Our approach obviates the need to determine outside dominator candidates for internal microtree vertices, eliminating the additional complexity Alstrup et al. require to manage this information. We can thus summarize the key dierences in the various approaches as follows. Harel [1985] and Alstrup et al. [1997] partition D into microtrees and apply the standard LT algorithm to all of D, using precomputed tables to speed the computation of the link-eval data structure in the microtrees. We also partition D into microtrees, but we apply the LT algorithm, with the link-eval data structure unchanged, only to one, big region of D and use memoization to speed the computation of dominators in the microtrees. In other words, Harel [1985] and Alstrup et al. [1997] take a purely data structures approach, leaving the LT algorithm unchanged but employing sophisticated new data structures to improve its running time. We modify the LT algorithm so that, although it becomes slightly more complicated, simple and standard data structures su-ce to implement it. A minor dierence in the two approaches regards the use of tables. Harel [1985] and Alstrup et al. [1997] precompute the answers to all possible queries on mi- crotrees and then use table lookup to answer the queries during the actual dominators computation. We build the corresponding table incrementally using memo- ization, computing only the entries actually needed by the given instance. The two approaches have identical asymptotic time complexities, but memoization tends to outperform a priori tabulation in practice, because the former does not compute 6 Adam L. Buchsbaum et al. answers to queries that will never be needed. 3.2 Microtrees Consider the following procedure that marks certain vertices in D. The parameter g is given, and initially all vertices are unmarked. For x in D in reverse DFS order do y child of x S(y) If S(x) > g then Mark all children of x endif done Mark root(D) For any vertex v, let nma(v) be the nearest marked ancestor (not necessarily proper) of v. The nma function partitions the vertices of D into microtrees as follows. Let v be a marked vertex; T vg is the microtree containing all vertices x such that v is the nearest marked ancestor of x. We say is the root of microtree T v . For any vertex x, micro(x) is the microtree containing x. See Figure 1. For any v, if v has more than g descendents, all children of v are marked. There- fore, each microtree has size at most g. We call a microtree nontrivial if it contains a leaf of D. Only nontrivial microtrees can contain more than one vertex; these are the subtrees we process using memoization. The remaining microtrees, which we call trivial, are each composed of singleton, internal vertices of D; these vertices comprise the upper subtree, D 0 , of D. Additionally, all the children of a vertex that forms a trivial microtree are themselves microtree roots. Call a vertex v that forms a trivial microtree special if each child of v is the root of a nontrivial microtree. (In Figure 1(b), c and e are special vertices.) If we were to remove the nontrivial microtrees from D, these special vertices would be the leaves of the resulting tree. Since each special vertex has more than g descendents, and the descendents of any two special vertices form disjoint sets, there are O(n=g) special vertices. We note that Alstrup et al. [1997] dene microtrees only where they include leaves of D (our nontrivial microtrees), whereas our denition makes every vertex a member of some microtree. We could adopt the Alstrup et al. [1997] denition, but dening a microtree for each vertex allows more uniformity in our discussion, particularly in the statements and proofs of our lemmas and theorems. Gabow and Tarjan [1985] pioneered the use of microtrees to produce a linear-time disjoint set union algorithm for the special case when the unions are known in advance. In that work, microtrees are combined into microsets, and precomputed tables are generated for the microsets. Dixon and Tarjan [1997] introduce the idea of processing microtrees only at the bottom of a tree. 3.3 Path Denitions v) be a path in G. We dene P to be an external dominator path (abbreviated xdom path) if P is an sdom path and dominator path is simply a semidominator A New, Simpler Linear-Time Dominators Algorithm 7 path that resides wholly outside the microtree of the target vertex (until it hits the target vertex). The external dominator of vertex v is there is an xdom path from u to vgg : In particular, for any vertex v that forms a singleton microtree, We dene P to be a pushed external dominator path (abbreviated pxdom path) nontrivial microtrees occur only at the bottom of D, a pxdom path to v cannot exit and reenter micro(v): to do so would require traversing a back arc to a proper ancestor of root(micro(v)). Therefore, a pxdom path to v is (a) an xdom path to some vertex x 2 micro(v) catenated with (b) an x-to-v path inside micro(v). Either (a) or (b) may be the null path. The pushed external dominator of vertex v is there is a pxdom path from u to vg: Note that pxdom(v) 62 micro(v): since the arc (p D (root(micro(v))); root(micro(v))) catenated with the tree path root(micro(v)) pxdom path to v, we have that pxdom(v) pD (root(micro(v))). For example, consider vertices l and h in DFS tree D in Figure 1(b). The DFS number of l is 4. The path is an sdom path from r to l, and so is not an xdom path. Path (c; j; l) is an xdom path from c to l, and no xdom path exists from r to l, so 2. P is a pxdom path, however: (r; b; e; n) is an xdom path from r to n 2 micro(l), and (n; l) is a path internal to micro(l). Thus, semi(l). Continuing, the DFS number of h is 12. The only sdom path to h is (g; h), and so Path (b; d; f; g; h) is a pxdom path to h, however, so In general, for any vertex, its semi, xdom, and pxdom values need not match. We use the following lemmas. Note the similarity of Lemma 3.2 to Lemma 2.2. Lemma 3.1. For any vertex v that forms a singleton microtree, semi(v). Proof. Let v) be a pxdom path from u to v. If v forms a singleton microtree, then and so, by denition of pxdom, x i v for 1 i < k. Without loss of generality, however, since u is the minimum vertex from which there is a pxdom path to v, we can assume that x i 6= v for 1 i < k. Therefore P is a semidominator path, so semi(v) u. Any semidominator path, however, is a pxdom path, so in fact Lemma 3.2. idom(v) 62 micro(v) =) idom(v) pxdom(v). Proof. Let As observed above, u 62 micro(v). By denition of pxdom, there is a path from u to v that avoids all vertices (other than u) on the tree path u Therefore, if idom(v) 62 micro(v), idom(v) cannot lie on that tree path. In the next section, we give the details of our algorithm. 4. DETAILS OF OUR LINEAR-TIME ALGORITHM At a high level, we can abstract our algorithm as follows: 8 Adam L. Buchsbaum et al. a d f (a) a d f (b) Fig. 2. (a) The microtree T , consisting of vertices a, d, and f from Figure 1(b), as well as incident arcs external to T . (b) The induced graph aug(T ). (1) Using memoization to reduce running time, determine for each vertex v if so, the actual value idom(v). (2) Use the LT algorithm to compute idoms for all v such that idom(v) 62 micro(v). The remainder of this section provides the details behind our approach. For clarity, we describe as separate phases the resolution of the idom(v) 2 micro(v) question, the computation of pxdom(v), and the overall algorithm to compute idom(v). We discuss in Section 7 how to unite these phases into one traversal of D. 4.1 Computing Internal Dominators We begin by showing how to determine whether idom(v) 2 micro(v) and, if it is, how to nd the actual value idom(v). For vertex v that comprises a singleton microtree, our decision is trivial: idom(v) 62 micro(v). For a nonsingleton microtree T , we dene the following augmented graph. Let G(T ) be the subgraph of G induced by vertices of T . Let aug(T ) be the graph G(T ) plus the following: (1) A vertex t, which we call the root of aug(T ), or root(aug(T )). (2) An arc (t; v) for each v 2 T such that there exists an arc (u; v) 2 A for some We call these blue arcs. Note that there is a blue arc (t; root(T )). Vertex t represents the contraction of ignoring all arcs that exit T . See Figure 2. We use the augmented graphs to capture the intuition that removing arcs that exit a microtree 2 does not change the dominator relationship. We dene the internal immediate dominator (iidom) of vertex x, iidom(x), to be the immediate dominator of x in aug(micro(x)). We show that if iidom(x) 2 conversely, that if Computing iidoms using memoization on aug(micro(v)) thus yields a fast procedure to deter- exits micro(u) if v 62 micro(u). A New, Simpler Linear-Time Dominators Algorithm 9 y x (a) y z x (b) Fig. 3. aug(micro(x)), plus incident external arcs/paths from G. Solid lines are arcs; dotted lines are paths. (a) The case t, and z < y. If z 62 micro(x), e.g., in the gure, then there is a path from t to x in aug(T ), using blue arc (t; v), that avoids y. If in the gure, then there is a path internal to micro(T ) from z to x that avoids y. Either case contradicts the assumption that y = iidom(x). (b) Similar case, but y < z. There is a path P in aug(T ) from y to x, avoiding z. P contains no blue arcs, so it is a path in G, contradicting that z = idom(x). mine whether or not idom(v) 2 micro(v) for any v. We give the details of the memoization procedure below. Lemma 4.1. iidom(x) 6= root(aug(micro(x))) =) Proof. Let idom(x) such that y 6= t and y 6= z. If z < y, then in the full graph G, there exists a path P from z to x that avoids y. We use P to demonstrate a path P 0 in aug(T ) from some z 0 2 ft; zg to x that avoids y, contradicting the assumption that v) be the last arc on P such that u 62 aug(T ). If there is no such arc then P yields an immediate contradiction. Otherwise, arc (u; v) induces blue arc (t; v) 2 aug(T ). This arc together with a subpath of P from v to x provides path P 0 . See Figure 3(a). On the other hand, if y < z, then there is a path P in aug(T ) from y to x that avoids z. By hypothesis, y 6= t, so P contains no blue arcs. (There are no arcs into t, and so t 62 P .) Therefore, P is also a path in G, contradicting that z = idom(x). Figure 3(b). Lemma 4.2. Proof. Let t. Then there is a path P in aug(T ) from t to x that avoids z. If P contains no blue arcs, then it is a path in the original graph, contradicting the claim that z = idom(x). If P contains blue arc (t; v) for some v, then in G there is an arc (u; v) for some u 62 T . The tree path root(G) catenated with the arc (u; v) and the subpath in P from v to x gives a path in G to x that avoids a z x Fig. 4. aug(micro(x)), plus incident external arcs/paths from G. Solid lines are arcs; dotted lines are paths. Case in which t. There is a path P in aug(T ) from t to x, avoiding z. If P contains no blue arcs (follows the path from a to b around z), this contradicts that z = idom(x), for P exists in G. If P contains blue arc (t; v), then there is a path in G to x, using the arc (u; v), where u 62 T . Again, P avoids z, contradicting z, again giving a contradiction. See Figure 4. We memoize the computation of iidom(v) as follows. The rst time we compute the internal immediate dominators for some augmented graph aug(T ), we store the results in a table, I , indexed by graph aug(T ) and vertex v. We encode aug(T ) by a bit string corresponding to its adjacency matrix represented in row-major order. To compute this bit string, we traverse aug(T ) in DFS order, assigning DFS value one to the root of aug(T ) and using the DFS values as vertex identiers; we refer to this as the canonical encoding of aug(T ). If a subsequent microtree T 0 has an augmented graph that is isomorphic to encodings will be identical, so we can simply look up the iidom values for aug(T 0 ) in table I . This obviates having to recompute the iidoms for aug(T 0 simply map the iidom values stored in table I , which are relative to the canonical encoding of aug(T 0 ), to the current instantiation of aug(T 0 ). (Vertex x in aug(T 0 ) corresponds to vertex x root(aug(T in the canonical encoding of aug(T 0 ).) 4.2 Computing Pushed External Dominators We now prove that the following procedure labels vertices with their pxdoms. As we will show, this process allows us to avoid performing links and evals within nontrivial microtrees. Initially, D, and we use a link-eval data structure with label(v) as the value for vertex v. As we will see, by Theorem 4.4, pxdom(v) when v becomes linked. The link-eval values are thus pxdoms. micro(v)g, the external neighbors of v, be the vertices outside micro(v) with arcs to v. The procedure processes the microtrees T A New, Simpler Linear-Time Dominators Algorithm 11 in reverse DFS order. (1) For 6 (c) label(v) min(fvg Lemma 4.3 proves that this labels v with xdom(v). (2) For (v) be the set of all vertices in T from which there is a path to consisting only of arcs in G(T ). Set label(v) min y2Y (v) flabel(y)g. We call this pushing to v. (Pushing can be done by computing the strongly connected components of G(T ) and processing them in topological order.) Theorem 4.4 proves that pushing labels v with pxdom(v). If T is a trivial microtree, then link(v). Due to the pushing in Step (2), pxdom values are nonincreasing along paths from the microtree root. This allows us to perform evals only on parents of microtree roots: the pxdom pushing eectively substitutes for the evals on vertices inside the microtrees. To prove that the above procedure correctly labels vertices in a microtree T , we assume by induction that the procedure has already labeled by their pxdoms all vertices in all trees preceding T in reverse DFS order. The base case is vacuously true. Lemma 4.3. After Step (1), Proof. Let We show that (1) label(x) w and (2) label(x) w. (1) Consider the xdom path P from w to x. Let y 62 micro(x) be the last vertex on P before x. Let z be the least vertex excluding w that P touches on the tree path w else P is not an xdom path. The prex P 0 of P from w to z is a semidominator path. Otherwise, there exists some u 6= w on P 0 such that u < z; by Lemma 2.1, P 0 contains a common ancestor of u and z, contradicting the assertion that z is the least vertex in P on the tree path y. Therefore, pxdom(z) semi(z) w. By induction, label(z) w. If z 2 micro(y), then label(z) got pushed to y, and thus label(y) w. (Note that in Step (1).) If z 62 micro(y), then C in Step (1) contains some value no greater than label(z), due to the previous links via Step (3). In either case the label considered for x via the (y; x) arc is no greater than label(z) w. Figure 5. (2) Consider arc (y; x) such that y 62 micro(x). Let so there is a pxdom path P from w 0 to y. P catenated with the arc (y; x) is an xdom path. Similarly, for there is a pxdom path P from w label(z) to z. P catenated with the tree path z y and arc (y; x) forms an xdom path from w 0 to x. In either case, Figure 5 again demonstrates the potential paths. Theorem 4.4. After Step (2), 12 Adam L. Buchsbaum et al. y x Fig. 5. Microtrees containing x and y, with incident external paths. Solid lines are arcs; dotted lines are paths. There is an sdom path from w to some z > nca(y; x); thus in the gure), then label(y) w. If z 62 micro(y) in the gure), then label(eval(root(micro(y)))) w. Proof. We argue analogously to the proof of Lemma 4.3. Let we show that (1) w is considered as a label for x via an internal pushing path and (2) for any w 0 so considered, there is a valid pxdom path from w 0 to x. (1) Consider the pxdom path P from w to x. Let v be the rst vertex on P inside During Step (2), w is pushed to x via the path from v to x. (2) Consider any w 0 pushed to x. w 0 is an xdom or pxdom for some vertex y 2 T . there is a valid pxdom path from w 0 to x. 4.3 Computing Dominators Using the information we computed in Sections 4.1 and 4.2, we now give an algorithm to compute immediate dominators. The algorithm proceeds like the LT algorithm; in fact, on the subtree of D induced by the trivial microtrees, it is exactly the LT algorithm. The algorithm relies on the following two lemmas: Lemma 4.5. For any v, there exists a w 2 micro(v) such that (3 (4 Proof. The proof proceeds as follows. We rst nd an appropriate vertex w on the tree path root(micro(v)) v. We show that A New, Simpler Linear-Time Dominators Algorithm 13 x (a) x (b) Fig. 6. The graph induced by micro(v), plus incident external arcs/paths from G. Solid lines are arcs; dotted lines are paths. be the pxdom path from x to v. w is the least vertex in P on the path from root(micro(v)) to v. (a) The prex P 0 of P from x to w includes only vertices that are greater than v (except w). (b) P 0 includes descendents of w that are less than v and so must take a back arc to w. In either case, P 0 is an sdom path from x to w, since w is the least vertex in P on the path from root(micro(v)) to v. argue that This resolves postulates (1){(3). Finally, we prove that idom(w) 62 micro(x), which implies postulate (4). consider the pxdom path P from x to v. Let w be the least vertex in P on the tree path root(micro(v)) v. We argue that the prex P 0 of P from x to w is a semidominator path. If not, then there is some vertex y 6= x on P 0 such that y < w. Since w v, it must be that y 2 micro(v); otherwise, y violates the pxdom path denition, since we only allow a y < v on P if y 2 micro(v). By Lemma 2.1, the subpath of P 0 from y to w contains a common ancestor z of y and w. Since y < w, it must be that z < w. As with y, it must also be that z 2 micro(v), or else z violates the pxdom path denition. This implies that z is on the tree path root(micro(v)) v, contradicting the assertion that w is the least such vertex on P . Therefore semi(w) x. See Figure 6. Now we argue that semi(w) x. If not, there is a semidominator path P from some y < x to w. P catenated with the tree path from w to v, however, forms a pxdom path from y to v, contradicting the assumption that Similarly, we argue that any semidominator path is also a pxdom path, pxdom(w) x. If there is a pxdom path P from some y < x to however, P catenated with the tree path from w to v is a pxdom path that contradicts the assumption that x. Thus we have shown that By denition of pxdom, pxdom(w) < root(micro(w)). Therefore, semi(w) implies that semi(w) 62 micro(w). By Lemma 2.2, therefore, idom(w) 62 micro(w), and thus by Lemma 4.1, Lemma 4.6. Let w; v be vertices in a microtree T such that 14 Adam L. Buchsbaum et al. y Fig. 7. The graph induced by micro(v), plus incident external arcs/paths from G. Dotted lines are paths. If idom(v) < idom(w), then there is an sdom path from some y < idom(w) to some x > idom(w) such that there is a tree path from x to v. If x lies on the tree path from idom(w) to w in the gure), however, this contradicts the denition of idom(w), and if x lies on the tree path from w to v in the gure), this contradicts that (3 Proof. Condition (3) and Lemma 4.2 imply that idom(v); idom(w) 62 T . In particular, idom(v) < w, so Lemma 2.3 implies that idom(v) idom(w). If idom(v) < idom(w), then there is a path P from idom(v) to v that avoids idom(w). must contain a semidominator subpath P 0 from some y < idom(w) to some x > idom(w) such that x v. x cannot lie on tree path idom(w) would contradict the denition of idom(w). x cannot lie on tree path w v, for this would imply pxdom(v) y < pxdom(w). (By Lemma 3.2, idom(w) pxdom(w).) So no such P 0 can exist. See Figure 7. Lemmas 4.5 and 4.6 imply the following, which is formalized in the proof of Theorem 4.7. Consider a path in a microtree, from root to leaf. The vertices on the path are partitioned by pxdom, with pxdom values monotonically nonincreasing. Each vertex w at the top of a partition is such that thermore, idom(w) 62 micro(w). For another vertex v in the same partition as w, either idom(v) is actually in the partition, or else outside the microtree. See Figure 8. That implies that our algorithm devolves into the LT algorithm on the upper subtree, D 0 , of D consisting of trivial microtrees. We can now compute immediate dominators by Algorithm IDOM, given in Figure 9. For each v 2 D, IDOM either computes idom(v) or determines a proper ancestor A New, Simpler Linear-Time Dominators Algorithm 15 x y z a c Fig. 8. A microtree with incident external paths. Dotted lines are paths. y, and All vertices on the tree path from w to p(v) have pxdom y; the path from x to b, a prex of which is an xdom path, does not aect the pxdom values on the w-p(v) part of the partition. The vertices in the partition need not share idoms, however. In this picture, y, but u of v such that description of the straightforward postprocessing phase that resolves the latter identities. IDOM uses a second link-eval data structure, with pxdom(v) as the value for vertex v; at the beginning of IDOM, no links have been done. Theorem 4.7. Algorithm IDOM correctly assigns immediate dominators. Proof. Lemma 4.1 shows that assigning idom(v) to be iidom(v) if iidom(v) 2 micro(v) is correct. Assume then that iidom(v) 62 micro(v), and thus idom(v) 62 micro(v) by Lemma 4.2. Consider the processing of vertex v in bucket(u). Assume rst that be the child of u on tree path u v. We claim that z is the vertex on tree path u 0 v with minimum semi and that Assuming that this claim is true, if Observe that for any w 2 micro(v) such that w pxdom(w) semi(w). Thus, if v, and the claim holds. On the other hand, if then is the vertex on the tree path pD (root(micro(v))) of minimum pxdom. The claim holds, since (1) pxdom(u 0 ) Consider the remaining case, when pxdom(v) 6= semi(v). Lemma 4.5 shows that there exists a w 2 micro(v) such that w root(aug(micro(v))), and so pxdom(w) and and w are both placed in the same bucket Algorithm IDOM For in reverse DFS order do Process(v) done For such that fug is a trivial microtree, in reverse DFS order do link(u) done Process(v) If iidom(v) 2 micro(v) then else add v to bucket(pxdom(v)) endif For do If z v else z eval(p D (root(micro(v)))) endif idom(v) u else endif done Fig. 9. Algorithm IDOM. by IDOM. Therefore, IDOM does compute the same value for idom(v) as for idom(w), and by the previous argument, it computes the correct value for idom(w). 5. ANALYSIS Here we analyze the running time of our algorithm. It should be clear that the generation of the initial DFS tree D and the division of D into microtrees can be performed in linear time, by the discussion in Section 3.2. 5.1 Computation of iidoms Recall the memoized computation of iidoms described in Section 4.1. So that all the iidom computations run in linear time overall, the augmented graphs must be small enough so that (1) a unique description of each possible graph aug(T ) can be computed in O(jaug(T )j) time and (2) all the immediate dominators for all possible augmented graphs are computable in linear time. (After computing immediate dominators for an augmented graph, future table lookups take constant time each.) We require a description of aug(T ) to t in one computer word, which we assume holds log n bits. Recall each microtree has no more than g vertices, for some parameter g. Thus, each augmented graph has no more than g+1 vertices. Without aecting the time bounds (we can use g 1 in place of g), we can assume that any aug(T ) has no more than g vertices. Therefore, aug(T ) has no more than g 2 arcs A New, Simpler Linear-Time Dominators Algorithm 17 and can be uniquely described by a string of at most g 2 bits. To t in one computer word, We can traverse aug(T ) and compute its bitstring identier in O(jaug(T )j) time, assuming that we can (1) initialize a computer word to 0, and (2) set a bit in a computer word, both in O(1) time. This further assumes that vertices in T are numbered from 1 to jT j, where jT j is the number of vertices in T . As part of the DFS of G, we can assign secondary DFS numbers to each v, relative to root(micro(v)), satisfying this labeling constraint. The total time to generate bitstring identiers is thus O microtree T Since each vertex (respectively, arc) in G can be attributed to one vertex (respec- tively, arc) in exactly one augmented graph, and there is one extra root vertex for each augmented graph, Expression When rst encountering a particular aug(T ), we can use any naive dominators algorithm to compute its immediate dominators in poly(g) time. Then we can store the values for iidom(v), for each v 2 aug(T ), in table I in time O(jaug(T )j). In the worst case, we would have to memoize all the iidom values for all possible distinct graphs on g or fewer vertices. There are about 2 g 2 such graphs, so the total time is O(2 poly(g)), inducing the constraint poly(g) n: A simple analysis shows that if using memoization, we can compute all needed iidom values in O(n +m) time. 5.2 Computation of pxdoms Step (1) in the computation, the initial labeling of a vertex v, processes each vertex and arc in G once throughout the labelings of all vertices v. Additionally, Step (1) performs at most one eval operation, on a trivial microtree root, per arc in G. Step (2) can be implemented by computing the strongly connected components (SCCs) of the subgraph of G induced by the microtree T , initially assigning each vertex in each SCC the minimum label among all the vertices in the SCC, and then pushing the labels through the SCCs in topological order. Computing SCCs can be done in linear time [Tarjan 1972], as can the topological processing of the SCCs. Step (3) links root(T ) once for each trivial microtree T . Thus, the time to compute the pxdoms, summed over all the microtrees, is n) plus the time to perform at most n 1 link and m eval operations. We analyze the link-eval time in Section 6. 5.3 Computation of idoms We implement the bucket associated with each vertex by a linked list. For each takes constant time to look up iidom(v) and either assign idom(v) or place v into bucket(pxdom(v)). To process a vertex v in bucket(pxdom(v)) requires constant time plus the time to perform eval on pD (root(micro(v))). Each vertex appears in at most one bucket, so processing the buckets takes time O(n) plus the time to do at most n evals on trivial microtree roots. (Since pxdom(v) 62 micro(v), only trivial microtree roots have buckets.) Again, we perform link(v) only on trivial microtree roots, so the total time taken by IDOM is O(m n) plus the link-eval time. 5.4 Summary By the above analysis, the total time required to compute immediate dominators in a owgraph G with n vertices and m arcs is O(m+n) plus the time to perform the links and evals on D. We next prove that since we do links and evals only on trivial microtree roots, the total link-eval time is O(m n) for an appropriate choice of the parameter g. 6. DISJOINT SET UNION WITH BOTTOM-UP LINKING Recall that link and eval are based on disjoint set union, yielding the (m; n) term in the LT time bound. Here we show that restricting the tree to which we apply links and evals to have few leaves results in the corresponding set union operations requiring only linear time. Let U be a set of n vertices, initially partitioned into singleton sets. The sets are subject to the standard disjoint set union operations. and C are the names of sets; the operation unites sets A and B and names the result C. nd(u). Returns the name of the set containing u. It is well known [Tarjan and van Leeuwen 1984] that n 1 unions intermixed with m nds can be performed in O(m(m;n) + n) time. The sets are represented by trees in a forest. A union operation links the root of one tree to the root of another. Operation nd(u) traces the path from u to the root of the tree containing u. By linking the smaller tree as a child of the root of the larger tree during a union and compressing the path from u to the root of the tree containing u during nd(u), the above time bound is achieved. We show that given su-cient restrictions on the order of the unions, we can improve the above time bound. We know of no previous result based on this type of restriction. Previously, Gabow and Tarjan [1985] used a priori knowledge of the unordered set of unions to implement the union and nd operations in O(m time. We do not require advance knowledge of the unions themselves, only that their order be constrained. Other results on improved bounds for path compression [Buchsbaum et al. 1995; Loebl and Nesetril 1997; Lucas 1990] generally restrict the order in which nds, not unions, are performed. Of the n vertices, designate l to be special and the remainder n l to be ordinary. The following theorem shows that by requiring the unions to \favor" a small set of vertices, the time bound becomes linear. Theorem 6.1. Consider n vertices such that l are special and the remaining n l are ordinary. Let be a sequence of n 1 unions and m nds such that each A New, Simpler Linear-Time Dominators Algorithm 19 union involves at least one set that contains at least one special vertex. Then the operations can be performed in O(m(m; l) Proof. The restriction on the unions ensures that at all times while the sequence is being processed, each set either contains at least one special vertex or is a singleton set containing an ordinary vertex. This observation can be proved by an induction on the number of unions. The following algorithm can be used to maintain the sets. A standard union-nd data structure is created containing all the special vertices as singleton sets. Recall that such a data structure consists of a forest of rooted trees built on the vertices, one tree per set. The root of a tree contains the name of the set. There is also an array, indexed by name, that maps a set name to the root of the corresponding tree. We will call this smaller data structure U 0 and denote unions and nds on it by union 0 and nd 0 . The ordinary vertices are kept separate. Each ordinary vertex contains a pointer that is initially null. The operations are performed as follows. If each of x and y names a set that contains at least one special vertex, perform union 0 (x; Suppose one of x and y, say y, is a singleton set containing an ordinary vertex. Set the pointer of the ordinary vertex to point to the root of set x. Relabel that root z. nd(x). If x is a special vertex, execute nd 0 (x). If x is ordinary and has a null pointer, return x. (It is in a singleton set.) If x is ordinary with a nonnull pointer to special vertex y, return nd 0 (y). The intuition is simple: unless an ordinary vertex x forms a singleton set, it can be equated to a special vertex y such that Each operation involves O(1) steps plus, possibly, an operation on a union-nd data structure U 0 containing l vertices. Let k be the total number of operations done on U 0 . Then the total running time is O(k(k; l)+m+n), which is O(m(m; l)+n) It is convenient to implement the above algorithm completely within the frame-work of a single standard union-nd forest data structure, using path compression and union by size, as follows. Initially all special vertices are given weight one, and all ordinary vertices are given weight zero. Recall that the size of a vertex is the sum of the weights of its descendents, including itself. To see that this implementation is essentially equivalent to that described in Theorem 6.1, observe the following points. First, by induction on the number of operations, an ordinary vertex is always a leaf in the union-nd forest. The union- by-size rule ensures that whenever a singleton ordinary set is united with a set containing special elements, the ordinary vertex is made a child of the root of the other set. The standard nd operation is done by following parent pointers to the root and then resetting all vertices on the path to point to the root. Hence any leaf vertex, and in particular any ordinary vertex, remains a leaf in the forest. Each ordinary vertex is thus either a singleton root or contains a pointer to a special vertex, as in the proof of Theorem 6.1. Furthermore, since the ordinary vertices have weight zero they do not aect the size decisions made when uniting sets containing special vertices. A nd on an ordinary vertex is equivalent to a 20 Adam L. Buchsbaum et al. nd on its parent, which is a special vertex, just as in the proof of Theorem 6.1. The only dierence is that the pointer in the ordinary vertex is possibly changed to point to a dierent special vertex, the root. This only adds O(1) to the running time. 6.1 Bottom-Up linking Let a sequence of unions on U be described by a rooted, undirected union tree, T , each vertex of which corresponds to an element of U . The edges in T are labeled zero or one; initially, they are all labeled zero. Vertices connected by a path in T of edges labeled one are in the same set. Labeling an edge fv; p(v)g one corresponds to uniting the sets containing v and p(v). The union sequence has the bottom-up linking property if no edge fv; p(v)g is labeled one until all edges in the subtree rooted at v are labeled one. Corollary 6.2. Let T be a union tree with l leaves and the bottom-up linking property. Then n 1 unions and m nds can be performed in O(m(m; l) time. Proof. Let the leaves of T be classed as special and all internal vertices classed as ordinary. When the union indicated by edge fx; p(x)g occurs, all descendants of x, and in particular at least one leaf, are in the same set as x. Therefore the union sequence has the property in the hypothesis of Theorem 6.1. Alstrup et al. [1997] prove a variant of Corollary 6.2, with the m(m; l) term replaced by (l log l +m), which su-ces for their purposes. They derive the weaker result by processing long paths of unary vertices in T outside the standard set union data structure. We apply the standard set union data structure directly to T ; we need only weight the leaves of T one and the internal vertices of T zero. 6.2 Application to Dominators Recall the denition of special vertices from Section 3.2: a vertex is special if all of its children are roots of nontrivial microtrees. Theorem 6.3. The (n) links and (m) evals performed during the computation of pxdoms and by the algorithm IDOM require O(n +m) time. Proof. Consider the subtree T of D induced by the trivial microtree roots. All the links and evals are performed on vertices of T . The special vertices of D are precisely the leaves of T . We can view T as the union tree induced by the links. The links are performed bottom-up, due to the reverse DFS processing order. As shown in Section 3.2, there are O(n=g) special vertices in D and thus O(n=g) leaves of T . We choose log 1=3 n, which su-ces to compute iidoms in linear time. By Corollary 6.2, the link-eval time is thus O(m(m; n= log 1=3 n) n). The theorem follows, since m n and (m; Our algorithm is completely general: it runs in linear time for any input owgraph G. Corollary 6.2, however, implies that, implementing union-nd as described above, the standard LT algorithm [Lengauer and Tarjan 1979] actually runs in linear time for all classes of graphs in which the corresponding DFS trees D have the following property: the number l of leaves of D is su-ciently sublinear in m so that (m; A New, Simpler Linear-Time Dominators Algorithm 21 7. IMPLEMENTATION This section describes our implementation, which diers somewhat from our earlier description of the algorithm for e-ciency reasons. The input is a owgraph in adjacency list format, i.e., each vertex v is associated with a list of its successors. Figure presents the top-level routines, which initialize the computation, perform a depth-rst search and partition the DFS tree into microtrees, and compute dominators. The initialization code creates and initializes the memoization tables. The partitioning code assigns DFS numbers, initializes the vertices and stores them in an array, vertices, in DFS order, computes the size of the subtree rooted at each vertex, and identies microtrees using the subtree sizes. Each vertex is marked Plain, MTRoot, or TrivMTRoot, depending on whether it is a nonroot vertex in a microtree, the root of a nontrivial microtree, or the root of a trivial microtree. Also, each vertex is assigned a weight to be used by the link-eval computation: special vertices (recall that a vertex is special if all of its children are roots of nontrivial microtrees) have weight one, and ordinary vertices have weight zero. (See Lengauer and Tarjan [1979] for the implementation of link and eval.) Finally, we initialize an array, pmtroot, to contain parent(v) for each v. This array will eventually store parent(root(micro(v))) for each v. By initializing it to the vertex parent, we only have to update it for vertices in nontrivial microtrees, which we will do in ProcessMT below. The code to compute dominators given the partitioned DFS tree diers from our earlier presentation in two ways. First, we combine the processing of vertices and buckets into a single pass, to eliminate a pass over the vertex set, as do Lengauer and Tarjan [1979]. Second, we separate the code for processing trivial microtrees from the code for processing nontrivial microtrees, which allows us to specialize the algorithm to each situation, resulting in simpler and more e-cient code. These changes, which are simple rearrangements of the code, do not alter the time complexity of the algorithm. Computedom calls ProcessV, to handle trivial microtrees, and ProcessMT, to handle nontrivial microtrees. ProcessV, shown in Figure 11, computes the xdom and pxdom of v, stores v in the appropriate bucket, links v to its parent, and then processes the bucket of v's parent. This code exhibits both of our changes. First, we follow the LT approach to combining the processing of vertices and buckets: we link v to p, its parent, and then process p's bucket. Immediately following the processing of v, only vertices from the subtree rooted at v are in p's bucket. Adding the link from v to p completes the path from any such vertex to p, which allows us to process the bucket. Second, we exploit that idom(v) is guaranteed to be outside v's microtree, thereby eliminating a conditional expression. ProcessMT (Figure 12) performs similar steps but is more complex, because it processes an entire microtree at once. The rst step is to nd the microtree's root. Since the vertices in a microtree have contiguous DFS numbers, we can nd the root by searching backward from v in the vertices array for the rst vertex that is marked as a nontrivial microtree root. Once we have the microtree root, we update pmtroot(v) appropriately for each v in the microtree. Then we (1) compute the xdom of each vertex in the microtree and an encoding for the augmented graph that corresponds to the microtree, (2) compute iidoms, (3) compute pxdoms, and 22 Adam L. Buchsbaum et al. Initialize computation Partition(root) status(root) TrivMTRoot Computedom(root) Partition(Vertex v) Assign DFS number to v Mark v as visited bucket(v) NULL link(v) NULL label(v) dfsnum(v) status(v) Plain For s 2 successors(v) do If s has not been visited then endif Add v to predecessors(s) done If size(v) > g then If all v's children are Plain then Mark the Plain children of v in DFS tree with MTRoot status(v) TrivMTRoot endif Computedom(Vertex root) For in reverse DFS order do ProcessV(v) elseif v has not been processed ProcessMT(v) endif done For in DFS order do If samedom(v) 6= NULL then endif done Fig. 10. Pseudocode for computing dominators. A New, Simpler Linear-Time Dominators Algorithm 23 ProcessV(Vertex v) label(v) dfsnum(v) For do If label(p) < label(v) then label(v) label(p) endif If dfsnum(p) > dfsnum(v) then evalnode If label(evalnode) < label(v) then label(v) label(evalnode) endif endif done Add v to bucket(vertices[label(v)]) For do z else samedom(w) z endif delete w from bucket(parent(v)) done Fig. 11. Pseudocode for processing trivial microtrees. (4) process the bucket of the parent of the microtree root. We compute xdoms and the microtree encoding together, because both computations examine predecessor arcs. The microtree encoding is simple: two bits for each pair of microtree vertices, plus one bit for each blue arc. During this computa- tion, we also identify a special class of microtrees: a microtree is isolated if the only target of a blue arc is the microtree root. We will use this information to speed the computation of pxdoms. The iidom computation uses memoization to maintain the linear time bound. To increase its eectiveness, we remove unnecessary bits and eliminate unnecessary information from the microtree encoding used to index the memoization tables. First, we remove the bits for self-loops. Second, we exploit that a blue arc to v implies that iidom(v) 62 micro(v) and that none of the information about v's internal arcs is useful. In particular, since we know that the root of the microtree is always the target of some blue arc, we eliminate from the encoding the bits for arcs into the root. These changes reduce the size of the iidom encoding from g 2 +g to bits (from 12 bits to six, for In addition to reducing the size of the encoding, we can reduce the number of populated slots in the memoization table, using the same observation. If there is a blue arc into a nonroot vertex, w, we zero the remaining bits for arcs into it, because they are irrelevant. We do not ProcessMT(Vertex v) Find mtroot in vertices starting from v initialize encoding isolated true For do label(v) dfsnum(v) For do Include (p; v) in encoding else Include blue arc to v in encoding If v 6= mtroot then isolated false endif If label(p) < label(v) then label(v) label(p) endif If dfsnum(p) > dfsnum(v) then evalnode If label(evalnode) < label(v) then label(v) label(evalnode) endif endif endif done done iidomencoding reduced encoding If iidommemo[iidomencoding] is not dened then iidommemo[iidomencoding] Computeiidom(encoding) endif iidoms iidommemo[iidomencoding] If (isolated) then else endif For do delete w from bucket(parent(mtroot)) done Fig. 12. Pseudocode for processing nontrivial microtrees. A New, Simpler Linear-Time Dominators Algorithm 25 IsolatedPush(microtree MT; int iidoms[ ]) mtroot MT [0] for do label(v) label(mtroot) done Add mtroot to bucket(label(mtroot)) Fig. 13. Pseudocode for pushing in isolated microtrees. remove these bits, because we want a xed-length encoding. (The bits are extra only when there is a blue arc into w.) Once we compute the reduced encoding, we look it up in the memoization table to determine if futher computation is necessary to determine the iidoms. We use the O(n 2 )-time bit-vector algorithm [Aho et al. 1986] augmented to exploit the blue arcs, when necessary. The iidoms are expressed in terms of a DFS numbering of the augmented graph. We translate the augmented graph vertex into the corresponding vertex in the current microtree by adding its secondary DFS number to the primary DFS number of the root of the microtree. We have implemented two forms of Push. The rst, shown in Figure 13, is a simplied form that can be used for isolated microtrees. The absence of blue arcs into nonroot vertices implies that (1) the xdom of the microtree's root is the pxdom of all vertices in the microtree and (2) the immediate dominators of nonroot vertices will be local to the microtree (that is, root(micro(v))). The root vertex has a nonlocal idom, so we simply add it to the bucket of its pxdom. The second, shown in Figure 14, handles the general case. First, we compute strongly connected components (SCCs) using memoization. In this case, the memoization is used only for e-ciency. As with the iidom calculation, we use a reduced encoding for SCCs. The SCC encoding, which uses g 2 g bits, does not include self-loops or blue arcs, since neither aects the computation. We compute SCCs using the linear-time two-pass algorithm from Cormen et al. [1991]. Given the SCCs (either from the memoization table or by computing them), we process them in topological order: we nd the minimum of the xdoms of the vertices within the SCC and the incoming pxdoms, and then assign this value to each vertex as its pxdom. Given v's pxdom and iidom, we either assign idom(v) directly, if or we put v into the appropriate bucket. After pushing, we nish by processing the bucket of pmt, the parent of the mi- crotree's root. Any vertex in the bucket must have pmt as its immediate dominator. By denition, the pxdom of a vertex, v, in the microtree, is the minimum pxdom along the path from pmt to v. As a result, we can skip the eval on any vertex in pmt's bucket and assign pmt as the immediate dominator directly. 8. RESULTS This section describes our experimental results. It would be interesting to compare our algorithm (BKRW) with that of Alstrup et al. (AHLT) [1997], to judge relative constant factors, but AHLT relies on the atomic heaps of Fredman and Willard. 26 Adam L. Buchsbaum et al. mtroot MT [0] sccencoding reduced encoding If sccmemo[sccencoding] is not dened then sccmemo[sccencoding] Computescc(mtroot; encoding) endif For in topological order do For do If iidoms[secdfsnum(v)] 62 MT then Add v to bucket(label(v)) else endif done done Fig. 14. Pseudocode for pushing in the general case. Atomic heaps, in turn, are composed of Q-heaps, which can store only log 1=4 n elements given O(n) preprocessing time. The atomic heap construction requires Q-heaps that store 12 log 1=5 n elements. For atomic heaps, and thus the AHLT algorithm, to run in linear time, therefore, n must exceed 2 12 20 [Fredman and (Alternatively, one can consider AHLT to run in linear time, but with an impractically high additive constant term.) Alstrup et al. [1997] provide variants of their algorithm that do not use atomic heaps, but none of these runs in linear time. Ours is thus the only implementable linear-time algorithm, and we therefore compare our implementation of BKRW with an implementation of the LT algorithm derived from their paper [Lengauer and Tarjan 1979]. We performed two sets of experiments. The rst set used owgraphs collected from the SPEC 95 benchmark suite [SPEC 1995], using the CFG library from the Machine SUIF compiler [Holloway and Young 1997] from Harvard. 3 (Six les from the integer suite could not be compiled by Machine SUIF v. 1.1.2 and are omitted from the data.) The second set used some large graphs collected from our Lab. We performed our experiments on one processor of an eight-processor SGI Origin 2000 with 2048MB of memory. Each processing node has an R10000 processor with 32KB data and instruction caches and a 4MB unied secondary cache. Both implementations were compiled with the Mongoose C compiler version 7.0. We report aggregate numbers for the SPEC test set, because it contains a large number of owgraphs. Table I reports the sizes of the owgraphs, averaged by benchmark. Table II contains average running times for LT and for BKRW with microtree sizes of two and three. Figure 15 displays a scatter plot in which each 3 Machine SUIF is an extension of the SUIF compiler [Amarasinghe et al. 1995] from Stanford. We used Machine SUIF version 1.1.2. A New, Simpler Linear-Time Dominators Algorithm 27 Table I. Graph Sizes, Averaged Over the Flowgraphs in Each Benchmark, for the SPEC 95 Flowgraphs Benchmark Number of Average Average Flowgraphs Vertices Arcs CINT95 Suite 129.compress 132.ijpeg 524 14 20 147.vortex 923 23 34 134.perl 215 CFP95 Suite 145.fpppp 37 19 26 103.su2cor 37 104.hydro2d 43 35 46 125.turb3d 24 52 71 Table II. Running Times on the SPEC 95 Flowgraphs, Averaged Over the Flow- graphs in Each Benchmark. The Numbers in Parentheses Measure the Dier- ence Between the Two Algorithms, as Computed by the Following Formula: LT 100:0. Positive Numbers Indicate That LT is Better; Negative Numbers Indicate That BKRW is Better. Benchmark LT BKRW CINT95 Suite 130.li 20.01 us 33.91 us (69.49%) 36.99 us (84.90%) 129.compress 22.61 us 37.62 us (66.42%) 43.84 us (93.92%) 132.ijpeg 25.46 us 40.43 us (58.78%) 45.86 us (80.11%) 147.vortex 36.70 us 53.59 us (46.02%) 61.29 us (67.00%) us 56.61 us (43.93%) 63.73 us (62.04%) 099.go 50.39 us 69.87 us (38.66%) 79.37 us (57.51%) 126.gcc 66.56 us 87.87 us (32.01%) 95.61 us (43.63%) 134.perl 89.54 us 112.23 us (25.34%) 121.13 us (35.28%) CFP95 Suite 145.fpppp 32.75 us 46.63 us (42.37%) 49.33 us (50.61%) us 53.14 us (37.93%) 59.90 us (55.47%) 107.mgrid 38.36 us 53.72 us (40.02%) 60.32 us (57.23%) 103.su2cor 46.74 us 62.06 us (32.78%) 67.28 us (43.95%) 104.hydro2d 49.99 us 66.71 us (33.45%) 72.82 us (45.68%) 146.wave5 51.71 us 68.45 us (32.38%) 74.46 us (44.00%) 125.turb3d 73.66 us 103.56 us (40.60%) 110.36 us (49.82%) us 99.01 us (26.60%) 106.72 us (36.46%) 101.tomcatv 174.60 us 210.20 us (20.39%) 215.20 us (23.25%) 28 Adam L. Buchsbaum et al. Number of vertices1.52.5 Number of vertices101000 Number of flowgraphs Fig. 15. Relative dierences in running times of BKRW and LT (for 3). There is a point in the top plot for each owgraph generated from the SPEC 95 benchmarks. The bottom plot displays the number of owgraphs for each respective number of vertices. (Note that the y-axis on the bottom plot represents a logarithmic scale.) point represents the running time of BKRW (with microtrees of size two or three) relative to LT on a single owgraph. The plot shows that the overhead of BKRW is larger than that of LT on small graphs, but that the dierence tails o to about 10% quickly. For this gure, we combined the data from the integer and oating-point suites; separating the two, as in Table II, would yield two similar plots. Table III lists our large test graphs, which come from a variety of sources, along A New, Simpler Linear-Time Dominators Algorithm 29 Table III. Graph Sizes for the Large Test Graphs Graph Vertices Arcs ATIS 4950 515080 Phone 2048 4095 7166 with their sizes. The ATIS, NAB, and PW graphs are derived from weighted nite-state automata used in automatic speech recognition [Pereira and Riley 1997; Pereira et al. 1994] by removing weights, labels, and multiple arcs. The Phone graph represents telephone calling patterns. The augmented binary graphs (AB1 and AB2) were generated synthetically by building a binary tree of a given size (shown in the table as the graph's label) and then replacing each leaf by a sub- graph. See Figure 16. The AB1 graphs use the subgraph shown in Figure 16(b), and the AB2 graphs use the subgraph shown in Figure 16(c). These graphs were designed to distinguish BKRW from LT. The subgraphs will be treated as isolated microtrees in BKRW, which means that all the nonroot vertices in a microtree will have dominators within the microtree and that the back and cross arcs will be handled cheaply (without evals) by BKRW. In particular, calls to eval related to these arcs will be avoided by BKRW, and as a result, no links in the link-eval forest will be compressed by BKRW. We observed that, as expected, BKRW performs fewer links and evals than does LT. Running time is a more telling metric, however, and we present the running times for our experiments in Table IV. For the speech and Phone graphs, the overhead of processing the microtrees, which includes initializing the memoization tables, computing iidoms, computing microtree encodings, and pushing, outweighs the savings on calls to link and eval. BKRW does outperform LT on the larger augmented binary graphs. This is to be expected, since BKRW has substantially fewer calls to eval and compresses zero links for these graphs. In addition, the overhead of processing the microtrees is low, because they are all isolated. Note that the improvement of BKRW over LT decreases as the graphs get larger. The benet gained by our algorithm is small relative to the cost due to paging, which increases as the graphs get larger. (a)1(b)13 (c) Fig. 16. (a) T k is a k-depth binary tree. Augmented binary graph AB1 (respectively, AB2) is generated by replacing each leaf M i with the subgraph shown in (b) (respectively, (c)). Table IV. Running Times on the Large Test Graphs. The Numbers in Parentheses Measure the Dierence Between the Two Algorithms, as Computed by the Following LT 100:0. Positive Numbers Indicate That LT is Better; Negative Numbers Indicate That BKRW is Better. Graph LT BKRW ATIS 384.50 ms 423.38 ms (10.11%) 427.25 ms (11.12%) ms 2836.25 ms (3.55%) 2844.75 ms (3.86%) ms 3195.98 ms (2.20%) 3189.15 ms (1.99%) Phone 8313.62 ms 8594.75 ms (3.38%) 8616.38 ms (3.64%) 1024 2.00 ms 2.00 ms (0.00%) 2.50 ms (25.00%) 2048 5.00 ms 5.00 ms (0.00%) 5.00 ms (0.00%) ms 10.00 ms (-9.09%) 12.00 ms (9.09%) ms 22.00 ms (-9.74%) 23.00 ms (-5.64%) ms 48.38 ms (-8.08%) 48.38 ms (-8.08%) ms 117.50 ms (-7.75%) 117.88 ms (-7.46%) 2097152 20188.00 ms 19499.00 ms (-3.41%) 19498.38 ms (-3.42%) 1024 4.00 ms 3.12 ms (-21.88%) 4.00 ms (0.00%) 2048 8.00 ms 7.00 ms (-12.50%) 7.00 ms (-12.50%) ms 15.62 ms (-2.34%) 16.00 ms (0.00%) ms 34.25 ms (-4.86%) 32.75 ms (-9.03%) ms 74.25 ms (-5.11%) 70.38 ms (-10.06%) ms 182.12 ms (-4.14%) 175.38 ms (-7.70%) 2097152 51920.62 ms 51461.25 ms (-0.88%) 51029.88 ms (-1.72%) A New, Simpler Linear-Time Dominators Algorithm 31 Given the overhead that BKRW pays for computing microtree encodings and pushing and that is very small, BKRW is surprisingly competitive, even for small owgraphs, but these experiments suggest that LT is the algorithm of choice for most current practical applications. LT is simpler than BKRW and performs better on most graphs. BKRW performs better only on graphs that have a high percentage of isolated microtrees. 9. CONCLUSION We have presented a new linear-time dominators algorithm that is simpler than previous such algorithms. We have implemented our algorithm, and experimental results show that the constant factors are low. Rather than decompose an entire graph into microtrees, as in Harel's approach to dominators, our path-compression result allows microtree processing to be restricted to the \bottom" of a tree traversal of the graph. We have applied this technique [Buchsbaum et al. 1998] to simplify previous linear-time algorithms for least common ancestors, minimum spanning tree (MST) verication, and randomized MST construction. We also show [Buchsbaum et al. 1998] how to apply our techniques on pointer machines [Tarjan 1979b], which allows them to be implemented in pure functional languages. ACKNOWLEDGMENTS We thank Bob Tarjan, Mikkel Thorup, and Phong Vo for helpful discussions, Glenn Holloway for his help with Machine SUIF, and James Abello for providing the phone graph. --R Compilers: Principles The Theory of Parsing Dominators in linear time. Manuscript available at ftp://ftp. The SUIF compiler for scalable parallel machines. Modern Compiler Implementation in C. Introduction to Algorithms. Optimal parallel veri The program dependency graph and its uses in optimization. A linear-time algorithm for a special case of disjoint set union A linear time algorithm for The ow analysis and transformation libraries of Machine SUIF. A fast algorithm for Linearity of strong postorder. Object code optimization. Postorder disjoint set union is linear. Speech recognition by composition of weighted In Finite-State Language Processing Weighted rational transductions and their application to human language processing. Optimal control dependence computation and the Roman chariots problem. Algorithm 430: Immediate predominators in a directed graph. Finding dominators in directed graphs. Applications of path compression on balanced trees. A class of algorithms which require nonlinear time to maintain disjoint sets. accepted June --TR Worst-case Analysis of Set Union Algorithms Compilers: principles, techniques, and tools A linear algorithm for finding dominators in flow graphs and related problems The program dependence graph and its use in optimization Introduction to algorithms Postorder disjoint set union is linear Efficiently computing static single assignment form and the control dependence graph Trans-dichotomous algorithms for minimum spanning trees and shortest paths Data-Structural Bootstrapping, Linear Path Compression, and Catenable Heap-Ordered Double-Ended Queues Linearity and unprovability of set union problem strategies I Optimal control dependence computation and the Roman chariots problem Modern compiler implementation in Java Linear-time pointer-machine algorithms for least common ancestors, MST verification, and dominators Applications of Path Compression on Balanced Trees A fast algorithm for finding dominators in a flowgraph Immediate predominators in a directed graph [H] Object code optimization The Theory of Parsing, Translation, and Compiling --CTR Adam L. Buchsbaum , Haim Kaplan , Anne Rogers , Jeffery R. Westbrook, Corrigendum: a new, simpler linear-time dominators algorithm, ACM Transactions on Programming Languages and Systems (TOPLAS), v.27 n.3, p.383-387, May 2005 G. Ramalingam, On loops, dominators, and dominance frontier, ACM SIGPLAN Notices, v.35 n.5, p.233-241, May 2000 On loops, dominators, and dominance frontiers, ACM Transactions on Programming Languages and Systems (TOPLAS), v.24 n.5, p.455-490, September 2002 Andrzej S. Murawski , C.-H. Luke Ong, Fast verification of MLL proof nets via IMLL, ACM Transactions on Computational Logic (TOCL), v.7 n.3, p.473-498, July 2006 Loukas Georgiadis , Robert E. Tarjan, Finding dominators revisited: extended abstract, Proceedings of the fifteenth annual ACM-SIAM symposium on Discrete algorithms, January 11-14, 2004, New Orleans, Louisiana Loukas Georgiadis , Robert E. Tarjan, Dominator tree verification and vertex-disjoint paths, Proceedings of the sixteenth annual ACM-SIAM symposium on Discrete algorithms, January 23-25, 2005, Vancouver, British Columbia Efficient Algorithm for Finding Double-Vertex Dominators in Circuit Graphs, Proceedings of the conference on Design, Automation and Test in Europe, p.406-411, March 07-11, 2005 Adam Buchsbaum , Yih-Farn Chen , Huale Huang , Eleftherios Koutsofios , John Mocenigo , Anne Rogers , Michael Jankowsky , Spiros Mancoridis, Visualizing and Analyzing Software Infrastructures, IEEE Software, v.18 n.5, p.62-70, September 2001 Ren Krenz , Elena Dubrova, A fast algorithm for finding common multiple-vertex dominators in circuit graphs, Proceedings of the 2005 conference on Asia South Pacific design automation, January 18-21, 2005, Shanghai, China Elena Dubrova, Structural Testing Based on Minimum Kernels, Proceedings of the conference on Design, Automation and Test in Europe, p.1168-1173, March 07-11, 2005

compilers;dominators;microtrees;path compression;flowgraphs

295681

Coordinating agent activities in knowledge discovery processes.

Knowledge discovery in databases (KDD) is an increasingly widespread activity. KDD processes may entail the use of a large number of data manipulation and analysis techniques, and new techniques are being developed on an ongoing basis. A challenge for the effective use of KDD is coordinating the use of these techniques, which may be highly specialized, conditional and contingent. Additionally, the understanding and validity of KDD results can depend critically on the processes by which they were derived. We propose to use process programming to address the coordination of agents in the use of KDD techniques. We illustrate this approach using the process language Little-JIL to program a representative bivariate regression process. With Little-JIL programs we can clearly capture the coordination of KDD activities, including control flow, pre- and post-requisites, exception handling, and resource usage.

INTRODUCTION KDD-knowledge discovery in databases-has become a widespread activity undertaken by an increasing number and variety of industrial, governmental, and research organizations. KDD is used to address diverse and often unprecedented questions on issues ranging from marketing, to fraud detection, to Web analysis, to command and control. To support these diverse needs, researchers have devised scores of techniques for data preparation, transformation, mining, and postprocessing. Moreover, dozens of new techniques are added each year. While the growing collection of techniques and tools helps address the growing set of needs, the size and rapid growth of the collection is becoming something of a problem itself. Many of the techniques will yield incorrect results unless they are used correctly with other techniques. In addition, KDD is often done by teams whose activities must be correctly coordinated. Thus, one of the chief challenges facing an organization that wishes to conduct KDD is in assuring that data analysis and processing techniques are used appropriately and correctly and that the activities of teams assembled to do KDD are properly controlled and coordinated. The applicability of techniques can depend on a number of factors, including the question to be addressed, the characteristics of the data being studied, and the history of processing of those data. This problem can be compounded if the organization lacks experience with the (possibly new) techniques, or if individual analysts on a team differ with respect to their general level of expertise, specialized knowledge about the data (e.g., biases and assumptions), or familiarity with particular analysis techniques (pitfalls and tricks). The problem can be further exacerbated if multiple analysts must be orchestrated in a KDD effort, or if the resources required to support the KDD effort are scarce or subject to competitive access. We view these problems as issues of coordination, with the general goal being to assure that the right team member applies the right technique to the right data at the right time. Similar problems of coordination come up in software development, for example, in the application of software tools to software artifacts, the assignment of developers to development tasks, and the organization of tasks in the execution of software methods. We have applied process programming to solve coordination problems in software development [17, 18], and we believe that process programming is also suited to representing and supporting coordination in KDD processes. The applicability of approaches based on software process programming is further suggested by other similarities between KDD processes and software processes. For example, both sorts of problems entail the involvement of both human and automated agents, the combination of algorithmic and non-algorithmic techniques, the reliance on external resources, and the need to react to contingencies and handle exceptions. Additionally, issues of process are important in understanding and assuring the validity of KDD results. In this paper we argue that a process orientation is important for KDD and that process programming is an appropriate technique for effecting good coordination in the Copyright - 1999 by the Association for Computing Machinery, Inc. Permission to make digital or hard copies of part or all of this work for personal or classroom use is granted without fee provided that copies are not made or distributed for profit or commercial advantage and that copies bear this notice and the full citation on the first page. Copyrights for components of this work owned by others than ACM must be honored. Abstracting with credit is permitted. To copy otherwise, to republish, to post on servers or to redistribtue to lists, requires prior specific permission and/or a fee. use of KDD techniques. We support this argument with examples programmed in Little-JIL, a process language that emphasizes coordination of activities, agents, and the use of resources and artifacts. We believe that Little-JIL provides a basis for orchestrating coordination that assures correctness and consistency in the specification and execution of KDD processes, and assures that agents will have the ability to communicate, analyze, and generally reason about the coordination of KDD techniques. 2. KDD PROCESSES A process can be thought of as a multi-step plan for completing a given task. A process specification defines a class of process instances. Each instance conforms to the specification, but carries out its work in ways that are molded by the mix of agents and data that are available when the process is executed. Instances differ from each other in ways that include the selection of agents that execute particular steps, the order in which steps are executed, and the choice of which substeps are used to complete a given step. For example, a single KDD process specification for bivariate regression might allow choice among multiple methods for handling outliers (e.g., manual removal, automatic removal, non-removal), for constructing a regression model (simple least-squares regression, locally- weighted regression, and three group resistant line), and for estimating statistical significance (parametric estimates, randomization tests). Naively assuming no interstep constraints and only these three steps, this very simple process can be instantiated in different ways - a potentially confusing number for an unaided user. Some of these possible configurations of process steps are clearly more desirable and effective than others in different situations. Thus researchers and practitioners have begun to provide this sort of guidance. Presently this takes the form of technical papers that specify desirable processes in informal ways. We believe that there is considerable value in augmenting these informal descriptions with the more precise, complete, and formal specifications that are achievable through process programming. Capturing and representing processes precisely, completely, and clearly is notoriously difficult, but our preliminary work indicates that carefully designed process specification languages can greatly facilitate this task. are Particularly Important to KDD Explicit representation of processes is particularly important in KDD. First, effective KDD requires managing dependencies between steps. Some steps may require, disallow, or enable other steps. For example, using most neural network training algorithms requires a preceding step to recode missing values. Non-parametric regression techniques disallow any subsequent step to construct parametric confidence intervals. Constructing a decision tree enables a future step of pruning that tree. Explicit representations of these dependencies can assure that they are appropriately handled. Second, the details of processes are essential to determining the statistical validity of inductive inferences. One example of this is the well-known error of testing on training data [24]. KDD processes that do not enforce separation between training and testing data (e.g., through simple disjoint sets or cross-validation) will produce biased estimates of model accuracy. The underlying cause of this phenomenon - referred to as "multiple comparisons" in statistics - has far more general effects. It has been causally linked to several pathologies of data mining algorithms, including attribute selection errors, overfitting, and oversearching [14] and pathological growth in the size of decision trees [15]. It has also been causally linked to errors in evaluating several types of modeling algorithms [8, 11, 12]. KDD systems that employ multiple analysts distributed in time and space are particularly susceptible to pathologies stemming from multiple comparisons [16]. Explicit representation of KDD processes supports analyses that can determine when these pathologies can and cannot occur. In addition, the ability to reinvoke an identical process is a necessary prerequisite to solutions such as randomization tests, cross-validation, and bootstrap estimates [20]. Explicit representation of processes provides a vehicle for assuring that reinvocations are indeed identical. Third, process details are vital to establishing the validity of KDD results in more general ways. The literature of KDD, statistics, and machine learning is filled with discoveries of implicit assumptions underlying particular techniques. In most cases, the only way to verify whether these assumptions are met is to examine the process used to apply a particular technique. Only by knowing the process used to derive a result can potential errors be traced back to their source. Explicit KDD process descriptions capture these details. Fourth, explicit representation of KDD processes can help balance multiple performance goals. Several approaches to a given analysis task may produce results of differing statistical validity, comprehensibility, and ultimate utility. In addition, those techniques may require different amounts of computation effort and human attention. By explicitly representing these characteristics as part of the specification of individual steps, the process specification can be created that meets particular objectives (e.g., "give me a fast approximate result" or "give me a highly accurate result, but take all night if you need it"). Combining Human Analysts and Automated Agen t s Research on KDD processes represents a return to one of the central issues of early work in KDD: how best to combine the goals and expertise of human users with powerful automated data analysis tools. While this topic was identified as a central one by early work in the field (e.g., [9]), it can be overlooked in our rush to develop more sophisticated automated techniques. Recent work has returned to this theme, including general descriptions of KDD processes (e.g., [7]), analysis and integration of steps [6, 26], formulation of exploratory data analysis as an AI planning activity [21], and a nascent industry effort to formulate standard KDD processes (CRISP-DM (see http://www.ncr.dk/CRISP/)). More broadly, we believe that the effective integration of the work of human and automated agents is a problem that is at the core of a growing number of critical problems. We believe that we can advance work on this problem by studying it in the more specific context of mixed-agent coordination in KDD process specification. One important note: our work explores how to coordinate the activities of multiple KDD agents, be they automated or human. Our work does not concern programming individual automated agents for such tasks as training a neural network or calculating a chi-square statistic. These tasks are best done using conventional programming languages and software engineering techniques. Our work also does not attempt to tell human analysts how to do their job. Human analysts have knowledge and expertise that is essential to the KDD process. Instead, we are exploring flexible languages that can be used to coordinate the actions of experienced human analysts with those of automated agents and to build processes that enable less experienced analysts to achieve high-quality results. The next section provides an extended example of one such language. 3. AN EXAMPLE: BIVARIATE REGRESSION In this section we present an example of a KDD process for bivariate regression. Regression appears to be a relatively simple process, but it is an appropriate example nevertheless. First, it is a common data analysis activity, regression tools are included in several KDD workbenches, and it is a basic task in deployed KDD applications. Second, the process is not actually as simple as it may appear. It involves a combination of human and automated agents, it may draw on a variety of analytical techniques, the use of these techniques may be conditional and contingent, interdependencies exist between certain techniques, and the whole process may entail sequential, parallel, alternative, and recursive activities. Thus, although bivariate regression is a relatively "small" process, it still suffers many of the coordination problems that process programming is intended to address. The basic bivariate regression problem can be described simply (see Figure 1a). We have a continuously-valued variable X (e.g., advertising spending), and we wish to determine whether it can help us predict another continuously-valued variable Y (e.g., net sales). To assess this relationship between X and Y, we have a data sample of N (x, y) tuples. In this section, we present a process that coordinates agents and techniques in the performance of bivariate regression. We begin with basic linear regression, and then expand the example to incorporate further functionality in the form of non-linear regression and accommodation of inhomogeneous data sets (i.e., data reflecting two or more independent phenomena). The process is defined using the process language [25], which is described with reference to the examples. This process should not be taken as a complete or comprehensive specification. It contains both intentional and unintentional simplifications. That said, we believe that it illustrates many of the necessary features of a more complete specification, and that the Little-JIL language could be used to represent many of the necessary details in a more complete specification. 3.1 Linear Regression The most common approach to the task of bivariate regression is linear regression. Linear regression constructs a model of the form assessment of the statistical significance of the slope - 1 . We can conclude that X and Y are dependent if we can reject the null hypothesis that - 1 is zero with high confidence. Least squares regression (LSR) is the most commonly used form of linear regression. The advantages of LSR include relatively high statistical power and computational efficiency. However, LSR's desirable characteristics rest on several assumptions, including homoskedasticity (the variance of Y is independent of X) and the absence of outliers- (x, y) tuples that lie far from all other points. Outliers often represent errors or highly unusual conditions that produce extreme values. An outlier Bivariate regression Inhomogeneity Y Figure 1: Simple bivariate regression and two common problems Consider the assumption about outliers in more detail. Outliers strongly affect LSR models-a single outlier can sharply shift an LSR model, causing it to accurately predict neither the outlier, nor the other data points (Figure 1b). An alternative modeling technique -three group regression [5]-is robust to the presence of outliers. TGR divides the range of X into three groups with equal numbers of points, finds the median X and Y value of each group, and constructs a line from those three points. Because the median is a measure of central tendency that is resistant to outliers, TGR is much less strongly affected by outliers than LSR. TGR addresses the problem of outliers, but the parametric significance test of - 1 used for LSR does not apply to TGR. Instead, a computationally-intensive-technique - randomization test [1, 4]- should be used to test significance for the slope of the line built with TGR. Incidentally, a randomization test can also be used for LSR (although, due to its computational cost, we chose to exclude this from our example process). How the varied activities of linear regression should be coordinated, in light of the relevant dependencies, conditions, alternatives, and contingencies, is precisely what a cogent process definition should make clear. Such process definitions require a process language that enables coordination semantics to be expressed clearly and concisely, that allows rigor and flexibility to be combined as appropriate, and that supports effective process enforcement while admitting dynamic adaptation. 3.2 Representing a Linear Regression Process In this section we illustrate the linear regression process using the Little-JIL process language. Little-JIL is a visual language derived from a subset of JIL, a process language originally developed for software development processes [22]. Little-JIL focuses on coordination of agents in the performance of process activities in a wide range of processes. Little-JIL represents the activities of a process as steps, where each step can be decomposed into substeps. Substeps within a step can be invoked either proactively or reactively. A step may also have a prerequisite to guard entry into the step, a postrequisite to guard exit from the step, and exception handlers to handle exceptions thrown during the step. The requisites and exception handlers in turn are steps that may also have substeps, etc. In addition, steps may include resource specifications. Runtime management of resource allocation provides another means of dynamically constraining, adapting, and controlling process execution. Each step also has, as a distinguished resource, an execution agent, which is responsible for initiating and carrying out the work of the step. Execution agents may be human or automated, and both types may be transparently combined in a Little-JIL process. These features and others are illustrated and discussed below with respect to the examples. Figure 2 shows a Little-JIL specification of a linear regression process. Process steps in Little-JIL are represented visually by a step name surrounded by several graphical badges that represent aspects of step semantics. The bar below the step represents control of substeps. The leftmost element in the control bar is a sequencing badge that indicates how substeps should be executed. For example, the Linear Regression step in Figure 2 contains a circle-with-slash badge that represents a "choice" control construct; this indicates that Linear Regression is executed by executing one of the alternatives Least Squares Regression or Three Group Regression. The agent, an analyst to whom the step is assigned, makes this choice. Least Squares Regression and Three Group Regression, in turn, are executed by executing a sequence of substeps, as indicated by the arrow control badge. (Two other proactive control badges, "try" and "parallel", are discussed with respect to later figures.) The rightmost element of a step control bar represents exception handlers. Exception handlers may be simple actions or more complex subprocesses, represented by Figure 1: Little-JIL specification for linear regression additional substeps. The simple actions include completing the step, continuing the step, restarting the step, and rethrowing the exception. In Figure 2, the exception handler for the Outliers exception (thrown by step Construct Linear Model) has no substep; rather, this handler simply traps the exception and continues the Linear Regression step, as indicated by the arrow badge associated with the exception handler. handler with a substep is shown in Figure 4.) In the context of a choice step, continuing after an exception means that the agent is offered a choice of the remaining alternatives. A step may also include reactions, which are attached as substeps to a badge in the center of the control bar (however, reactions are omitted here for the sake of simplicity). In the visual representation of Little-JIL steps, a circular badge above a step name represents the interface to the step. The interface includes resources needed by the step, as well as parameters sent into and out of the step, local data, and events and exceptions that may be thrown by the step. Execution agents are represented as a type of resource. Each step has an execution agent; if none is specified for a step, the execution agent is inherited from the step's parent. In Figures 2 and 3 the agents include both humans and automated tools. Data sets can also be modeled as resources. Several steps in the example throw exceptions (designated in the interface by an X). While much of the data flow between steps is shown in a simplified form, most of the data declarations have been omitted from the interfaces in the figures for the sake of brevity. A Little-JIL step may also have a prerequisite and/or a postrequisite. A prerequisite is indicated by a downward- pointing triangle on the left of the step name and a postrequisite is indicated by an upward-pointing triangle on the right. An empty triangle indicates no requisite; a filled triangle with text indicates the name of the specified requisite. The body of the requisite is a separately specified step (not shown in our figures) possibly containing multiple substeps. A requisite is successful if it terminates normally; if it fails, it throws an exception. For example, the step Construct LSR Model has the postrequisite No Outliers. If outliers exist, then the postrequisite throws the Outliers exception, which causes Construct LSR Model to fail. The parent step Least Squares Regression propagates the exception, which is handled by its parent Linear Regression. Clearly, there are many ways to add to the process specified in Figure 2. Additional pre- and post- requisites could be added to the LSR and TGR steps, data preprocessing steps could be added to improve the robustness of the process, and other approaches to regression could be added. The next section discusses one of the most important elaborations to the process: how to deal with non-linearity. 3.3 Coping with Non-linearity A common diagnostic technique for any form of linear regression is to examine a plot of residuals. Ideally, the residuals-the errors in Y left unexplained by a model-should not vary with X. A non-linear relationship between X and the residuals indicates a non-linear relationship between X and Y, one that is not adequately captured by the linear model. Checking for linear residuals can be represented in Little-JIL as a postrequisite for the Linear Regression step. What if this postrequisite fails? One solution would be to try a non-linear modeling technique such as locally-weighted regression or lowess [2]. Figure 3 shows a process that includes both the original Figure 2: Regression with substeps for linear and non-linear regression linear regression step and a new step for non-linear regression. The "try" sequencing badge on the root regression step indicates that non-linear regression is invoked only if linear regression fails. Given the current specification of linear regression, the principal reason the step might fail is the presence of non-linear residuals. Linear regression and non-linear regression are partitioned as separate alternatives because different processes are required to determine if linear and non-linear models indicate a relationship between X and Y. Linear regression tests a relatively simple statistical hypothesis (- regression relies on a step Evaluate Relationship in which a human analyst makes a qualitative judgement. To assist in that judgment, a step to construct confidence intervals has been added to non-linear regression, although analysts should be cautious to distinguish between confidence intervals and significance tests [1]. Note that the overall Regression process coordinates the work of human and non-human agents who participate at various levels in the process. As with linear regression, many additions to the regression process are possible. These include additional approaches to non-linear regression, more quantitative substitutes for the evaluate relationship step, and prerequisites for the regression step. The next section describes one particularly important prerequisite for regression-homogeneity. 3.4 Coping with Inhomogeneity A frequently overlooked assumption of regression is that the data sample is homogeneous-that it represents a single uniform phenomenon rather than two or more phenomena with fundamentally different behavior (Figure 1c). For example, inhomogeneity can occur when men and women have different physiological responses to some phenomenon, yet data from men and women are mixed together for purposes of analysis. In contrast to outliers, which often represent errors that cannot be explicitly modeled, inhomogeneity represents two or more distinct data regimes that require independent modeling. Figure 4 shows a Model Relationships process that handles inhomogeneous data. The process first attempts to apply regression testing to a given bivariate data set. However, the regression step is guarded by a prerequisite that tests the homogeneity of the data. This prerequisite assures that a single regression is not performed on heterogeneous data. If the prerequisite is violated, the exception NonHomogeneity is thrown, which is caught by an exception handler for Model Relationships. The recursive process Model Subsets handles the exception. At the top level Model Subsets is a sequence. The first substep, Choose a Subset, chooses a data subset from the inhomogeneous data set. The second substep is a parallel step, Use and Choose Next. This substep, in parallel, applies regression to the selected subset and recursively calls Model Subsets on the remaining part of the data set. By this recursion, Model Subsets iteratively models subsets of the original data set, completing normally when no more subsets are available (as indicated by the "check" badge on the exception handler for the exception NoSubsetAvailable). Figure 3: Handling inhomogeneous data By combining the parallel step with recursion, multiple data subsets may be modeled concurrently. Note that, in this formulation of the process, a chosen data subset is not guaranteed to be homogeneous. In that case, when the process Regression is called on the subset, the homogeneity prerequisite will again throw the NonHomogeneity exception, which will take control back to the exception handler for inhomogeneous data (i.e., Model Subsets). As an alternative, we could have put a test for homogeneity as a postrequisite on the Choose a Subset step. 4 . Coordinating Agents at Process Execution Time In the preceding sections we have shown how Little-JIL can be used to flexibly specify a process that manages inter-step process dependencies for multiple execution agents. In this section, we describe how the activities of these agents are coordinated when a process is instantiated and executed. The vehicle for agent coordination during process execution is an agenda management system (AMS). An agenda management system is a software system that is based on the metaphor of using agendas, or to-do lists, to coordinate the activities of various human and automated agents. In such a system, task execution assignments are made by placing agenda items on an agenda that is monitored by one or more execution agents. Different types of agenda items may be used to represent different kinds of tasks that an agent is asked to perform. Our agenda management system [19] is composed of a substrate that provides global access to AMS data, a set of root object types (agendas, agenda items, etc.), application-specific object types that extend the root types, and application-specific agent interfaces (e.g., GUIs for human agents). We have designed and implemented an AMS specifically to support the execution of Little-JIL processes. This AMS has five types of agenda items: one item type corresponds to each of the four Little-JIL step kinds, and one item type corresponds to a process step at its lowest level of decomposition. Each Little-JIL agenda item has many attributes, including step name, execution agent, current status, log, step instance parameters, throwable exceptions, and interpreter. The last attribute is provided because, as we illustrate below, the Little-JIL interpretation architecture allows each step to have its own interpreter instance. When a step of a process program is first instantiated, an agenda item of the appropriate type is created and its attribute values are set accordingly (e.g., status is set to "Posted," input parameters are given the correct values). As the process executes, the attribute values change accordingly (e.g., the execution agent sets output parameter values, status is changed). Thus, process program execution state is stored within the AMS. This approach to storing process state is similar to that used in the ProcessWall [13]. An agent typically monitors one or more agendas to receive tasks to perform. Multiple agendas are used because an agent may frequently be involved in several disjoint processes (or acting in roles that are logically disjoint). When an item is posted to an agenda that an agent is monitoring, the agent is notified that the agenda has changed. In the case of a human agent, for example, this could result in a new item appearing in the person's agenda view window. The agent is then responsible for interpreting the item and performing the appropriate task. Agents may also monitor items individually; this gives them the ability to post an item to an agenda and to observe the item so they can react to changes in the item's status, for example. These mechanics are sufficient for the Little-JIL interpreter to instantiate and execute multi-agent Little-JIL process programs. By examining the state of an agenda item corresponding to a step of the process program, the interpreter can execute the process. When a new step is to be executed, the interpreter identifies the appropriate execution agent (with the help of a resource management system), creates an appropriately typed agenda item for that step, and posts it on the agent's agenda. As the agent executes the step, its updated status is reflected in the agenda item's status attribute value, which is monitored by the interpreter. As the status changes, the interpreter accordingly creates and posts substeps, returns output parameters on successful completion of the step, propagates exceptions on unsuccessful completion, and so on. Thus, an AMS provides language-independent facilities that allow coordination to take place, while the interpreter encodes key coordination semantics of the Little-JIL language itself. This design decouples concerns about why and when coordination should occur from concerns about how coordination should occur. For example, consider how the process program fragment in Figure 2 would be executed, supposing an interpreter had created an item to correspond to an instance of a Linear Regression step (a Little-JIL choice step). Assume the interpreter has identified a HumanAnalyst for this task (named Herman), posted it to Herman's agenda, and started the item's interpreter (which is stored in the interpreter attribute of the item). At this point the human analyst would use the GUI to change the status attribute of the Linear Regression item to "Starting." Its interpreter would be notified of this change and would create two new agenda items that correspond to the Least Squares Regression and the Three Group Regression sub-steps, then set the parent item's status to ``Started.'' Because new agents are not specified for these steps, the interpreter would post them to Herman's agenda and would start interpreters for the new items. Herman's agenda GUI would render the agenda to clearly depict the subitems of the choice item as alternatives. Suppose Herman then chooses to start the Least Squares Regression step, changing its status to "Starting." At this point both the Linear Regression item's interpreter and the Least Squares Regression item's interpreter would be notified of the change. The Linear Regression item's interpreter would react by setting the status of the other sub-item (Three Group Regression) to "Retracted." This would cause the item to disappear from Herman's agenda. Meanwhile, the Least Squares Regression item's interpreter would create a new item for the first substep, Construct Model. Because the process specifies an LSRTool for that step, the new item would be posted to a particular LSRTool's agenda, then the Least Squares Regression item's status would be set to ``Started.'' Whatever agent was monitoring the LSRTool's agenda would then be notified that the tool's agenda has changed. This agent would extract whatever information was needed by the tool from the agenda item, set the item's status to ``Starting,'' and would invoke the LSRTool agent. Because Construct LSR Model is a leaf step, the item's interpreter immediately changes its state to "Started." When the LSRTool finishes, the tool's agent would set the status of the leaf step appropriately ("Completing" if successful or "Terminating" if not), and that step's interpreter would complete the transition of the leaf step to a final state. The interpreter for the Least Squares Regression item would be notified that the step has changed, and, depending on its status, would start the next sequential sub-step or would terminate the parent. As previously mentioned, Little-JIL makes no distinction between human and tool agents. Similarly, neither does the AMS. As seen in the above example, different agents interact with the AMS, and consequently with the running Little-JIL process, via customized agent interfaces. For humans, this interface may be a GUI that is used to change an item's status, signal exceptions, change parameters, etc. For COTS tools (such as the LSRTool, perhaps), this interface may be a wrapper agent that integrates the tool with the AMS, spawning the tool to perform tasks in response to agenda items being posted to the tool's agenda and reporting the results of tool execution by setting agenda item attributes (e.g., parameters, status) as required. Our early experiences support our belief that an agenda management system provides an appropriate metaphor for coordinating interaction in mixed-agent processes such as KDD. We intend to continue experimenting with the use of Little-JIL and the AMS to facilitate coordination in such processes. 5. LESSONS LEARNED Our experience using Little-JIL to specify KDD processes has been instructive. Many coordination aspects of KDD processes (including examples not described here) have been easily expressed using Little-JIL. For example, one aspect well handled by Little-JIL is the highly variable control requirements of KDD processes. Conversely, KDD processes have drawn on the full range of Little-JIL control constructs. In some cases, processes require extremely strict control, and Little-JIL allows us to indicate this (e.g., by executing substeps in a specified order). In other cases, only very loose control is needed, and the language allowed us to specify this as well (e.g., by allowing user choice or parallel execution). We believe that successful process languages for KDD must allow flexibility to program processes both strictly and loosely. Little-JIL's pre- and postrequisites are essential to effective coordination in processes. Prerequisites make explicit the assumptions that underlie a sampling or analysis technique; postrequisites make explicit the acceptance criteria for the successful completion of a technique. The ability to make assumptions and acceptance criteria explicit is important for making a process understandable, evaluating its correctness, assuring its consistent execution, and validating its results. Similarly, the ability to represent exceptions and exception handling is critical for process robustness, reliability, and safety. In our KDD examples, exception management is also crucial in specifying process control structures. While many descriptions of KDD techniques use nearly ideal data, most practitioners who attempt to apply these techniques quickly uncover hidden assumptions, leading to exceptions in idealized process models. The ability to indicate possible exceptions, specify how they are to be handled, and direct subsequent execution, is essential to coordinating KDD efforts in real-world applications. Resource management provides another dimension of coordination in processes. Flexibility in agent coordination is afforded because Little-JIL process can be written independently of the specific execution agents to which they will be bound at run time. Additionally, the control model of the language, in conjunction with the agenda management system, allows processes to be written transparently with respect to the issue of human versus automated agents. However, runtime allocation of agents allows dynamic orchestration of agent activities and enables the dynamic adaptation of process behaviors to agent availability. Similar degrees of flexibility and opportunities for dynamic control apply to resources in general. 6. FUTURE WORK Our work to date with Little-JIL has convinced us of the general utility of process specification. However, at least three important areas of work remain. First, additional experience with specifying processes is needed. We intend to increase the level of sophistication of the existing processes and to develop processes in other application domains. In particular, we have begun development of processes in the areas of coordination of robot and processes used in electronic commerce. Second, while we believe that Little-JIL specifications are easy to read and write compared to more algorithmic languages, we would like the KDD process to be extended by non-programmers. We imagine providing a more sophisticated process editor that would assist a KDD researcher by assisting with the insertion of appropriate steps with the necessary prerequisites, postrequisites, data flow and exception handling. Finally, the Little-JIL language itself is still under development and there are a number of issues we intend to address. We are investigating integrating an AI planner [10] and resource-based scheduler [23] with Little-JIL. Such mechanisms would allow us to schedule agents and other resources based on cost, availability for a specific time and duration, and expected quality of their results. The results from planning would help guide agents in their decision making at choice and parallel steps by identifying which substeps are most likely to satisfy the time, cost, and quality constraints for process instances. We are also investigating the use of static analysis techniques [3] on Little-JIL processes. Specifically, we wish to prove properties of Little-JIL processes such as ordering rules (Step A always executes before Step B) and non-local dependencies (if Step A is performed, Step B is eventually performed). There are also some extensions to Little-JIL itself that we want to consider further. In particular, KDD processes appear to need a more explicit means of representing non-local control flow dependencies. For example, in our example regression process, a parametric significance test is only applicable when least squares regression is used. Currently, this control flow dependency is captured via data flow in Little-JIL. That is, least squares regression results in the computation of intermediate values that are used in the parametric test. The precondition test for the parametric test checks whether those data exist, and, if they do not, prevents the parametric test from being used. A more direct means of expressing this control flow dependency would be preferable to hiding it within the data flow, as is currently done. A more direct means would also enable our static analysis techniques to reason more effectively about the behavior of the program. 7. CONCLUSIONS Knowledge discovery research is developing and exploiting a diverse and expanding set of data manipulation and analysis techniques. Not all analysts, or even all organizations, can have a thorough knowledge of how to correctly and effectively combine and deploy these techniques. Process programming provides an effective means for specifying the coordinated use of KDD techniques by agents in potentially complex KDD processes. As demonstrated in this paper, KDD process specifications written in Little-JIL express requirements on individual techniques and capture dependencies among techniques. Little-JIL is a high-level process language designed to support the specification of coordination in processes; offers appropriate control flow constructs, pre- and post-requisites, reactions, exception handling, agent specifications, and dynamic resource bindings. Little-JIL enables explicit representation of KDD processes, allows reasoning about those processes, and supports correct execution of the processes. In turn, this enables KDD applications to produce reliable and repeatable results, which is necessary for the effective use of data mining across a wide range of organizations. ACKNOWLEDGMENTS This work was supported in part by the Air Force Materiel Command, Air Force Research Laboratory, and the Defense Advanced Research Projects Agency under Contracts F30602-94-C-0137, F30602-97-2-0032, and F30602-93-C- 0100. --R Empirical Methods for Artificial Intelligence. Smoothing by local regression: principles and methods (with discussion). Data Flow Analysis for Verifying Properties of Concurrent Programs Randomization Tests. Resistant lines for x versus y. A guided tour through the data mining jungle From data mining to knowledge discovery in databases. Which method learns the most from data? A Negotiation-Based Interface between a Real-Time Scheduler and a Decision-Maker Statistical significance in inductive learning Presenting and analyzing the results of AI experiments: Data averaging and data snooping The ProcessWall: A Process State Server Approach to Process Programming Multiple comparisons in induction algorithms. Building Simple Models: A Case Study with Decision Trees. Unique Challenges of Managing Inductive Knowledge An Adaptable Generation Approach to Agenda Management Computer Intensive Methods for Testing Hypotheses: An Introduction. Complex Goal Criteria and its Application in Design-to-Criteria Scheduling Computer Systems That Learn. KDD process planning --TR Randomization tests The ProcessWall Statistical significance in inductive learning Data flow analysis for verifying properties of concurrent programs Empirical methods for artificial intelligence The design of a next-generation process language An adaptable generation approach to agenda management Multiple Comparisons in Induction Algorithms Programming Process Coordination in Little-JIL Building Simple Models Enhancing Design Methods to Support Real Design Processes

knowledge discovery process;process programming;agent coordination;knowledge representation;agenda management

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The design of an interactive online help desk in the Alexandria Digital Library.

In large software systems such as digital libraries, electronic commerce applications, and customer support systems, the user interface and system are often complex and difficult to navigate. It is necessary to provide users with interactive online support to help users learn how to effectively use these applications. Such online help facilities can include providing tutorials and animated demonstrations, synchronized activities between users and system supporting staff for real time instruction and guidance, multimedia communication with support staff such as chat, voice, and shared whiteboards, and tools for quick identification of user problems. In this paper, we investigate how such interactive online help support can be developed and provided in the context of a working system, the Alexandria Digital Library (ADL) for geospatially-referenced data. We developed an online help system, AlexHelp!. AlexHelp! supports collaborative sessions between the user and the librarian (support staff) that include activities such as map browsing and region selection, recorded demonstration sessions for the user, primitive tools for analyzing user sessions, and channels for voice and text based communications. The design of AlexHelp! is based on user activity logs, and the system is a light-weight software component that can be easily integrated into the ADL user interface client. A prototype of AlexHelp! is developed and integrated into the ADL client; both the ADL client and AlexHelp! are written in Java.

INTRODUCTION Online customer service systems such as "Call Centers" or "Customer Care Centers" have been widely used, e.g., home- banking, telephone registration, online shopping, airline ticket booking, and digital libraries, etc. However, most of Supported in part by NSF grants IRI-9411330 and IRI-9700370. To appear in Proc. ACM International Joint Conference on Work Activities Coordination and Collaboration, WACC '99, Feb. 22-25, San Francisco, CA 1999 those systems lack sufficient capabilities of interactive communication mechanisms necessary for providing online customers with more sophisticated support and help. The emergence of the World Wide Web provides some new options for the user support and help problem because multi-media information such as images, audio/video, and animation can be easily presented in addition to traditional text. On the other hand, because of this and rapid advances in technol- ogy, the software systems for existing applications are becoming more and more complicated and at the same time new applications are being quickly developed. Such applications include distance learning, computer-based training, electronic commerce applications, and more. We believe that a good package of online help facilities can not only compensate for deficiencies of user interfaces, but also will make very complex services easier to learn and use. At the core of the help facilities lies the cooperating software systems or modules enabling communication between the user and service (human) agents. The communication may be based on paradigms of specific software systems, in addition to the ASCII and voice channels. For example, in the context of a digital library for georeferenced information both the user and the information specialist at the help desk (at a distant site) should be able to have a view of the same interface session so that the user can see what the information specialist does exactly. Similar situations can also occur in sessions for seeking technical support for a software systems, making online reservations and orders, etc. In this paper, we present the design and implementation of an online interactive help system in the context of a digital library for geospatially referenced information. In particular, we report our experiences in the development and integration of an online help desk system Alex- Help!, for the Alexandria Digital Library (ADL) [FFL In ADL, the user starts a session by initiating an HTTP-based connection to the ADL server. During one session, the ADL client software has a relatively complex user interface that allows users to browse maps, zoom in/out, construct queries (searches), and manipulate the results of queries. The potential users of ADL do not necessarily have much experience in accessing computers, and dealing with spatial information; their knowledge about software systems may be very limited. Our objective is to provide "just in time" online (collaborative) help facilities for the users of the ADL system. The AlexHelp! system provides the following functional- ity. The user can request to establish a connection to a help desk. Once the connection is made, both the user and the information specialist will share a common view of the user's ADL session. The information specialist is now able to perform various actions on the ADL client interface, including those mentioned above, and all actions happen simultaneously at both sides. In addition to such synchronized session support, AlexHelp! also allows the presentation of pre-recorded demonstration sessions to the user, and showing or replaying user sessions to the help desk. AlexHelp! is designed using user activity logs (i.e., session logs), and is a light-weight software component that can be easily integrated into or layered on top of the ADL user interface client. Our design is different from many other collaborative systems that are developed from scratch with collaborative environment tools. This is because the ADL system is already operational, and using tools like Habanero would require rewriting much of the ADL client code and could significantly increase the cost of communication. Moreover, the advantage of our approach is that it saves time by not repeating the work already done. As pointed out in [BM94], a straightforward approach to constructing collaborative environments that combines different communication mechanisms together will not necessarily result in a good collaborative environment. Bergmann and Mudge found in their experiment that the successful use of their system requires much logistics and support. How- ever, what we find is that for a certain class of collaborative applications, careful design makes it possible to significantly reduce such costs by automating most repetitive tasks. Most of our work on AlexHelp! is focused on the ADL client. Generally speaking, AlexHelp! is quite different from typical groupware [Gru94b]. The collaborative model that AlexHelp! supports is much simpler, since collaboration only occurs between users of the system and the help desk. In a collaborative session, one participant is designated the master and the others are slaves. The master controls all the slaves' views of the client interface (essentially). During the session, either the user or the help desk can be the master and the master "token" can be passed among participants. Such transferring must be under the control of the help desk. The collaborative actions supported during the session are not very complex, and the nature of the connection between master and slaves is almost stateless. But clearly such an application in the context of ADL is also very representative for various other online assistance services. In the CSCW research community, several typologies have been proposed. In terms of the Grudin's typology [Gru94a], CSCW models are categorized in two dimensions, space and time, as follows. Different but Different and Same Predictable Unpredictable Same Meeting Work Shifts Team Rooms Facilitation L Different but Tele-, Video-, Electronic Collaborative A Predictable Desktop Mail Writing Conferencing Different but Interactive Computer Workflow Unpredictable Multicast Boards Seminar The AlexHelp! system supports a combination of these mod- els: in the space dimension, the collaboration can occur at different locations, either predictable or unpredictable; in the time dimension, it can be at the same time or at a different but predictable time. Our work is also related to synchronized web browsers such as [WR94, FLF94]. However, AlexHelp! is based on and integrated into the ADL client software rather than relying on web browsers; the design issues and implementation techniques are thus quite different. This paper is organized as follows. Section 2 presents the motivation of AlexHelp!. Section 3 outlines the overall design and system architecture. Section 4 focuses on the session log file, the key component used in AlexHelp!. Section 5 briefly summarizes and discusses some implementation is- sues. Conclusions and future work are included in Section 6. In this section, we first give a motivating example in the context of the Alexandria Digital Library system. We then discuss the general concept of online interactive help systems or help wizards and their applications. The Alexandria Digital Library (ADL) [FFL + 95] is a digital library for geospatially referenced data including maps, images, and text. ADL's graphical user interface is a client application implemented in Java; the user runs it locally (on the user's workstation), connecting to the ADL via the In- ternet. The ADL client includes the following components (see Figure 1). 1. Map Browser Window (window 1 in Figure 1). The Map Browser window allows users to interactively pan and zoom a two-dimensional map of the globe to locate their area(s) of interest. In addition, the user can select one or more areas on the map that may be used to constrain queries. The map is also used to display the spatial extent of the items retrieved from a query. 2. Search Window (window 2 in Figure 1). The Search window allows users to set query parameters; parameters include choosing the collection(s) to search, location (i.e., coordinates from the Map Browser), type of information (maps or aerial photographs from the the catalog or hydrologic features from the gazetteer), date options. 3. Workspace Window (window 3 in Figure 1). The Workspace window displays and allows manipulation of query results. A query history is also maintained here, allowing the ADL client to be returned to the state of a prior query. A Scan Display or metadata browser (bottom part of the workspace window) displays brief object metadata. Brief metadata includes title, format, access information, spatial/temporal references, and a small scale version of the original image if available. From this window, full metadata and access information can be requested. 4. Help Window (window 4 in Figure 1). The ADL client's Help Window is a typical example of today's software help systems. It shows, depending on what component in the graphical user interface is se- lected, help topics specific to that component. ADL has a wide range of potential users, many of whom may not have experience interacting with complex software sys- tems. The Help Window is of some assistance. A tutorial and walkthrough are also made available through the ADL homepage 1 . Clearly something more is required, however, since it is still difficult for inexperienced users Figure 1: The ADL client/User Interface to know how to use the system for their particular purposes Consider the following example. Suppose M is a student majoring in political science. M is working on a research paper about the civil population in the Santa Barbara area, and one of her friends suggests that she can get some useful data from ADL. She follows her friend's suggestion and launches an ADL client. Unfortunately, she quickly finds herself frustrated. She tries the help system, but is still confused and unable to learn how to effectively utilize ADL. She needs a quick answer. If she were in the university library, she could ask one of the librarians to help her find the information she's looking for. With AlexHelp!'s extension to the ADL help system, she has more choices than blindly searching for answers from the static help system. The following are possible alternatives. - She could try the online tutorial (Demo Sessions). She could download some pre-recorded sessions and replay them (the Demo Session player is discussed in Section 4). These sessions could demonstrate to her the basic operation of the ADL client. - If the online tutorial doesn't help, she could try the Help Wizard. She would be guided through a process where she answers questions based on the nature of her prob- lem. The Help Wizard would then direct her to several Demo Sessions that would hopefully help solve her problem - If the first two methods don't solve her problem, she could connect to the help desk at ADL for an interactive session. M and the help desk could communicate through text-based chat, phone, or online audio channels while the information specialist guides M through the process of constructing a query and evaluating the information returned from the query. M would actually see, on her own ADL client, exactly what she would need to do because the help desk controls M 's client while they talk. Clearly, these help desk services can help the student in this example. It is conceivable different kinds of users may prefer some of the service to others. To provide the help desk services in the above example, the following capabilities need to be developed. ffl Collaborative sessions By collaborative session, we mean that the ADL clients on both sides (the user and the help desk) are "con- nected"; that is, one client is in control of the session and sends its actions to the other client, which mirrors them on its graphical user interface. The clients can switch roles during the session, allowing the user and help desk to participate in a rich exchange of information. ffl Session replay (demo sessions) Support staff are not always immediately available. Instead of forcing users to wait until they can contact an information specialist, AlexHelp! uses the concept of demo sessions. Demo sessions are examples of using the ADL client that may be replayed on the client itself, showing the user successful ways to use the client. To support this concept, we developed a Session Player. This gives the user VCR-style control over a session; the user may pause, rewind, start and stop sessions. Different sessions show the user how to perform tasks of varying complexity, and for those that are more com- plex, it is very useful to be able to pause a session and replay subtle or difficult portions. In addition, the Session Player allows users to record their own sessions, which may be sent to support staff for offline analysis or detection of usage patterns. ffl Multimedia communication Interaction between the user and the help desk should not be restricted to manipulating or observing the ADL client. It seems reasonable to assume that a network connection of some type has been established between the help desk's and user's ADL clients; it could be utilized for more than just synchronizing the clients. Collaborative help systems should provide additional methods for com- text-based chat, voice, and even video could be used to allow participants to communicate. Shared whiteboards fit the paradigm as well, but in this case the "whiteboard" is the application itself - the reason the participants are connected is to explain the use of the application. In addition to the above capabilities, it is also very useful in such a context to provide: ffl Call waiting and forwarding In order to serve the users fairly and efficiently, our design calls for a way to queue help requests when there is more than one user that wants to participate in a collaborative session. Call waiting means that when the user requests to connect with the help desk for a collaborative session, if the help desk is not immediately available, the user is notified that there will be some period of waiting, and information such as expected wait time and the reasons for the current delay might be made available to the user. Another method to support interactive help is call for- warding: if the current information specialist is not proficient in the particular area in which the user needs help, the user's connection may be ``task switched'' to another information specialist that has the proper expertise. ffl Multicast collaborative sessions With multicast support, more than one slave mode client can be synchronized with a single master mode client - a desirable feature to support scheduled instructional/ demonstration sessions or distance learning. Based on the analysis of the requirements for online assistance in ADL and the design of the ADL client, we decided that the interactive online help facilities in AlexHelp! should be designed as an independent component that is easily integrated with the ADL client. Even though the AlexHelp! system is relatively simple and does not include everything listed at the end of Section 2, it succeeds as a good example of adding collaborative functions on top of the single-user version of a software package. The feedback from the development and user evaluation groups within the ADL project shows that our system provides a feasible and efficient way to support interactive online help in ADL. In this section we discuss the main issues in developing the AlexHelp! system, including functionality, design, and architecture. 3.1 Functionality Our prototype AlexHelp! provides the following facilities: ffl Collaborative session establishment and operation A simple dialog window is used to initiate the connection between two ADL clients. Once the connection is made, AlexHelp! runs in one of two modes: Slave (receive) mode In this mode, the ADL client uses AlexHelp! to listen for incoming messages from the other client; when a message (a remote event) is received, the remote graphical user interface event is "replayed" or duplicated on the Slave client. The user of the Slave client is unable to change the client's state. Master (send) mode In this mode, the AlexHelp! system sends local graphical user interface events to the Slave ADL client, where they are replayed. The state of the Master ADL client is thus reflected in the Slave ADL client. In particular, the ADL client Map Browser window is able to replay operations such as "Zoom In", "Zoom Out", "Pan", etc. ffl Session replay We also developed a mechanism for replaying recorded sessions. As we discuss below and in the next section, every ADL client records its operations in a log file. Alex- Help! utilizes the logs in both Collaborative Sessions and Replay. 3.2 Design Consideration of an appropriate collaborative model for the AlexHelp! system was driven in part by our goal of quickly developing a prototype that could be used to demonstrate the possibilities of extending and developing online help sys- tems. While we had access to the ADL client's code base, we also had the constraint that we could not change the basic functionality of the client; in other words, the client had to perform in all other respects exactly as it had before. Thus we were faced with adding multi-user ability to an existing single-user application that was not originally designed with multi-user access in mind, a task Grudin recommends as a reasonable way to approach building groupware [Gru94b]. Not being able to fundamentally alter the client prevented us from using frameworks or toolkits for building collaborative environments like NCSA's Habanero [CGJ or Sun's JSDT (Java Shared Data Toolkit) [JSDT]. Using a framework such as JSDT would entail in essence rebuilding the application. On the other hand, including the ADL client in an application that comprises a collaborative environment like Habanero would require the user to run the ADL client as a "component application" from within the new environment, making the ability to participate in a collaborative help session dependent on this new environment. This was deemed unacceptable for two reasons. First, the ADL client is meant to be a stand-alone application; it is the library user's interface to the ADL. Second, the ADL client has a fairly small footprint in terms of memory and network resources; including it in a larger framework would impose additional resource requirements. Our specification calls for a model of collaboration best described as turn taking or token passing: semantically, only one of the user or the help desk should have control of both clients at a given time. During the period of collaboration, one of the clients is in the "slave" or receive mode, and the other is in the "master" or send mode. Considering this simple model and our desired features within the context of a collaborative help system for ADL, there is no need to use complex concurrency control or shared object models [GM94, MR91]. The user interface of the client, say A, that is currently in receive mode is simply "locked;" it may receive and interpret events or messages from the other client, say B (i.e., the remote client), but the user is unable to change client A's state. When the clients trade roles, client A takes charge and the events generated at A are sent and duplicated at client B. We realized early on that the ADL client log facilities already in place were rich enough to duplicate the actions of one client at another client separated by space and/or time. This not only simplified design of the client-to-client commu- nication, but also makes it easy to generalize the client-to- client model into a one-to-many "multicast" model. Rather than capturing and packaging system or user interface events for transmission, which is a potentially complex endeavor even in a syntactically sweet language like Java [JAVA], the log entries can simply be sent to the remote client as they are generated. The use of session logs in AlexHelp! is discussed in detail in Section 4. 3.3 Architecture Initially, it was intended that the help desk and the ADL user would run the same program (on their respective work- stations). It was thought that the symmetry of both participants using the same application (i.e., one set of code) was good design. As the project progressed, it was decided that the help desk should be responsible for controlling a help ses- sion, and the necessary features for control should be built into the help desk's client. During a help session, the information specialist may turn control over to the remote user (i.e., the help desk's client becomes the slave); the information specialist should have the ability to recover control of the session if desired. However, because AlexHelp! will be rewritten along with the next version of the ADL client, development of a single application continued. The help desk's session management features were not implemented in our current prototype. AlexHelp! consists of three layers: the ADL Client layer, the Event Handler layer, and the Communication layer (Fig- ure 2). These three layers are described below. 3.3.1 ADL Client The ADL client is the primary interface to the ADL for users of the library. In its first incarnation, the ADL client was a Java applet suitable for running within a Java-capable browser such as Netscape's Navigator. The current version is a stand-alone application, built entirely in Java. Subsequent development is intended to produce both stand-alone and applet versions. As mentioned in Section 3.2, we were not at liberty to fundamentally change the ADL client's design or function- ality. We chose therefore to layer the additional functionality we required on top of the ADL client (see Figure 2). Note that in our prototype, the ADL client communicates directly with both the Event Handler layer and the Communication Figure 2 reflects our design rather than our current implementation. In an iterative software development model, we would choose in our next iteration (post prototype) to restrict the ADL client to interacting with a single layer, the Event Handler layer, in order to keep the module dependencies clean. 3.3.2 AlexHelp! Event Handler Layer The Event Handler is layered directly on top of the ADL client. Its primary function is to receive incoming remote events from the Communication Layer, and replay or reproduce them on the local ADL client. We noted in Section 3.2 that the existing logging mechanism in the ADL client is rich enough in content to allow us to duplicate or replay remote events on a local ADL client. This allowed us a very simple design for the Event Handler layer: when a remote event is received in the form of a log entry (a string of ASCII characters), it is parsed to determine what event should be triggered locally in the ADL client. The Event Handler then directs the local ADL client to perform the event. For example, if the remote event is a pan (horizontal or vertical movement) in the remote ADL client's Map Browser window, the Event Handler notifies the local ADL client, providing the direction to pan (north, south, east, west), and also the distance. This design for handling remote events is easily generalized for a distance learning or "multicasting" scenario: a single information specialist or instructor runs an ADL client in master mode, and the client's events are broadcast to a group of ADL client users, all of whose clients are operating in the slave mode. Each ADL client in slave mode receives and handles remote events in the manner described above. 3.3.3 AlexHelp! Communication Layer In the discussion of the collaborative model chosen for the AlexHelp! system we mentioned JSDT, an example of a framework for building collaborative applications that provides abstractions for typical components of such systems such as "channel," "client," and "server." Had we the opportunity of building AlexHelp! (and also the ADL client) from scratch, it is likely that we'd have chosen such a framework to ease the typically troublesome task of properly designing and building a system in which distributed communication is central to operation. However, it was decided that given the time and limited flexibility in terms of altering the ADL client, it would be easier and faster to use simple TCP/IP (sockets) to implement client-to-client communication. The choice was made easy because of Java's inclusion of networking abstractions as part of its core libraries [JAVA]. The java.net.* library provides objects that encapsulate network connections that are either connection-oriented or connectionless, and also provides a model for sending the same message to multiple recipients (useful for our distance-learning "multicast" model). The current implementation of the AlexHelp! system provides client-to-client connections using connection-oriented Java sockets and a turn-taking slave/master model; the next version of AlexHelp! will include the distance learning mode as well. AlexHelp!'s architecture is structured in layers so that it would be relatively easy to replace one of the layers with an alternative implementation - hopefully, it would be easy to the point that replacing one layer would require no modifications to the adjacent layers. The Communication Layer is meant to be no exception. In addition, we foresee utilizing the Communication Layer to apply additional communication mechanisms: text-based chat, voice, video, etc. ADL Client Event Handler Communication Layer Communication Layer ADL Client Event Handler ADL Client Layer Communication Layer ADL Server Event Layer Handler Figure 2: The Architecture of AlexHelp! In this section we discuss the central technique of using session logs in developing synchronized sessions in AlexHelp!. We also illustrate that session logs facilitate the design of session replay and give a general discussion on applying this technique in other contexts. Several of our early design meetings included personnel from the ADL development team. Among the many invaluable things we learned from them was the fact that they built a simple user activity log system into the client. As part of the event handling in the ADL client, certain user- and system-generated events are logged into an ASCII text file, which may then be analyzed after the fact. The original intent of the ADL development team was to use the activity logs for analysis of user activity and usage pattern discovery. We saw that if the logs were augmented to include more information about each event, we could then use the log entries to implement coordination of two physically separate ADL clients. Further, the logs could be viewed as a persistence mechanism, giving us a method for "session replay". We felt that this (persistence of an entire user's session) would be crucial to integrating AlexHelp! into the ADL's existing help system. Our experience developing AlexHelp! using session logs as both a way to forward local events to a remote client and as a persistence mechanism showed us that for a certain class of applications, log files can be instrumental in rapidly and simply augmenting existing help systems (or building new ones), and also adding multi-user capability to single-user applications. In Section 4.1, we discuss the format of the ADL client's log files, and explain some example log entries. In Section 4.2, we discuss the use of logs in controlling collaborative sessions in AlexHelp!. Section 4.3 examines how log files are used to implement VCR-style control of Demo Sessions (log replay). Section 4.4 discusses general application of of the log-based technique. 4.1 The ADL Client Log File The ADL client log consists of event summaries, each of which makes up a single log entry. The order in which entries appear in the log maps loosely to the order in which they occur; there may be variation between two logs recording identical sessions due to timing differences that can be tracked to the effects of multithreading; i.e., a system event may occur before a user-generated event, but the user event is logged first due to scheduling. We did not experience unexpected behavior related to such scheduling discrepancies. In addition to information recording an event (see examples below), a log entry also includes information identifying the particular session and client that generated the event. In a log entry, the information of interest (information that allows an event to be reconstructed from the log entry) actually makes up very little of the entry. In many cases, less than about one third of a log entry is of interest. For brevity, the other information (such as session or client iden- tification) has been removed from the log entries used in the discussion below. The following are examples of the ADL client's log en- tries, and comments related to using the logs with AlexHelp!. ffl client action - Map Mode: ZOOM IN This log entry is among the simplest types. It is generated when the user clicks with the mouse on the button labeled "ZOOM IN" on the Map Browser window. Similar log entries are made when the user clicks on the buttons labeled "ZOOM OUT," "SELECT," and "ERASE." ffl client action - New Extent: -205.3 -40.56 -25.29 49.44 This log entry occurs when the user clicks the mouse in the map on the Map Browser window when the selected mode is either ZOOM IN or ZOOM OUT. "New Extent" identifies this as a map resizing event. The parameters correspond to the lower left and upper right corners of a rectangular geographic region which defines the new extent of the map in the Map Browser window. The first two numbers are the latitude and longitude of the lower left corner; the second two are the latitude and longitude of the upper right corner. ffl client action - Query Region(s) Modified: This log entry occurs when the selected mode is SELECT and the user clicks and drags the mouse in the map on the Map Browser window. This action creates a selection box, which defines a rectangular geographic region, as in the previous log entry. The parameters are again the latitudes and longitudes of the lower left and upper right corners of the geospatial region, which may be used to constrain queries. Selection boxes may overlap each other. ffl client action - Query Region(s) Modified: In this log entry, a selection box is again drawn; in this case, however, several selection boxes already existed. Log entries for modifications to Query Regions (selection boxes), like the log entry above, simply enumerate all of the selection boxes, whether the change was adding or removing a selection box. The log entry above shows the five currently existing selection boxes. ffl client action - Map: Button Pressed: west This log entry shows a "pan" event (horizontally or vertically repositioning of the map in the Map Browser win- dow). 4.2 Logs in Collaborative/Synchronized Sessions The ADL client's log facility was central to our design and implementation of synchronized help sessions. Our specification calls for the ability to capture, distribute, and reproduce graphical user interface events. Capturing such events as they occur can be troublesome at best, and at worst im- possible. Although Java aids the programmer in this type of endeavor (see Section 3.2), it was clear to us that in light of the fact that the session log mechanism was already in place, utilizing it would permit implementation to proceed much faster than dealing with (intercepting) events at the graphical user interface level. Augmenting the ADL client in order to send log entries to a remote client was a simple matter of hooking into the module responsible for writing log entries to a file. Acting on the log entries that have been received from a remote client entails using the ADL client's public application programming interface (API) to set its internal state, which triggers any corresponding changes in the graphical user interface. As an example, consider again two ADL clients, A and B, connected in a collaborative session. Client A is being operated by the Alexandria help desk, and B is being operated by a user of the library. A is currently in the master (send) mode; B is in the slave (receive) mode. When the information specialist draws a selection box on client A's Map Browser, the log module in client A appends to the log file an entry corresponding to drawing the selection box. Since A is in the master mode, the log entry is also handed up to AlexHelp!'s Event Handler Layer. There, the log entry is forwarded to AlexHelp!'s Communication Layer, where it is sent over its network connection to client B. At client B, the log entry is received at the Communication Layer, which immediately hands it down to the Event Handler Layer. The Event Handler parses the log entry to determine if it is of interest; in this case, it finds that the log entry consists of parameters making up a list of selection boxes. The Event Handler uses ADL client B's public API to set its collection of selection boxes; client B then redraws its display in the Map Browser window to reflect the change. When events are received at an ADL client in the slave (receive) mode, the graphical user interface gives the user some kind of visual cues to alert him or her that something has changed: when the state of the Selection Box/Map Zoom mode buttons are changed, for example, the button that is now selected flashes red and white several times. Given a priori knowledge of the possible forms log entries can take, the task of parsing a log entry and deciding what to do in terms of reproducing the event is rather simple. Consequently, the Event Handler Layer, possibly the most complex module in AlexHelp!, is essentially a parser for a rather restricted language (see example log entries in Section 4.1). 4.3 VCR-Style Control of Log Replay (Demo Sessions) As in AlexHelp!'s collaborative sessions, the ADL client's session logs were important in the design and development of AlexHelp!'s Demo Sessions. Giving a user the ability to play and replay a series of graphical user interface events as if the user were operating a VCR is a powerful learning tool that goes beyond help systems that are text-based or rely on short animations. Using AlexHelp!'s model for Demo Sessions in conjunction with detailed explanations (easily supplied along with the log files as text or audio files), help systems can be developed that show usage of complex software systems in a way that users would otherwise only experience by using the system, perhaps in a trial-and-error fashion. Demo Sessions in AlexHelp! use log files in the same way that client B in Section 4.2 uses incoming log entries. An ADL client playing a Demo Session uses AlexHelp!'s Event Handler Layer to read a log file, parsing each log entry and carrying out the event described in the entry. AlexHelp! uses an additional window in the graphical user interface to give the user control over replay. The user may set the rate of playback, start and pause playback, single-step the session (i.e., play a single event and pause), rewind a single event, or rewind to the start of the session. 4.4 The Use and Utility of Log Files We envision (for ADL as well as other complex software systems) help systems that utilize log files in concert with a "Help Wizard"; the Help Wizard would present the user with a series of questions that help to focus the user towards one or several log files that show examples of sessions that closely resemble what the user will need to do to achieve his or her goals with the application. This is in stark contrast to many help systems in current applications, which seem to consist mainly of duplicating the functional descriptions of each user interface component in the help files. Help systems built with session logs can be viewed as an extension of support staff; even if the support staff and the end users operate in different time zones, making synchronized activities difficult to schedule, custom session logs can be created and forwarded to the user in lieu of an online meeting. For example, one of our recommendations for application of AlexHelp! in ADL is to maintain a page on the Alexandria World Wide Web site dedicated to listing demo sessions, from which users may download demo sessions showing queries similar to those the users hope to carry out and thus learn how to use the ADL interface from the example sessions. Users can, in turn, record their own session log for support staff analysis; what better summary of a problem using an application than a complete reproduction of the user's activity? 5 IMPLEMENTATION OF A RAPID PROTOTYPE In this section we briefly discuss issues related to the implementation of our current prototype. 5.1 Module Architecture The relationship of the AlexHelp! modules with the original ADL system is shown in Figure 3. The implementation details of the extension modules will be discussed in the following sections. Of interest in Figure 3 are the various working modes of the extended ADL client: 1. Stand-alone Mode This is the same as the sole working mode supported by the original ADL system. Native user interface (UI) events generated by the local windowing environment are sent to the User Event Handler and then get recorded locally by the event logging module. 2. Master Mode When the client is in the master mode, the native UI events are still processed by the User Event Handler. However, now they are also sent to the remote client via the network interface. 3. Slave Mode When the client is in slave mode, the user interface should not respond to local UI events but remote events instead. Remote events go through the Event Parser and then the UI highlighting module, which presents visual cues to the user that remote events are occurring. 4. VCR Mode When replaying session logs (retrieved from the network or locally), the Event Parser passes each event to the UI highlighting module, which passes the events to the User Event Handler after performing the necessary highlighting 5.2 Network Interface The network sublayer is encapsulated in an object which is responsible for maintaining the TCP link and some state information related to the connection. The reason for choosing (as opposed to UDP) is that a reliable link that guarantees delivery of synchronization messages is desired. A reliable link greatly simplifies the mode switching proto- col. In the current prototype, we have not added the mode switching control to the help desk's ADL client (since we decided to keep only one code base). When implemented, the information specialist will control mode switches. The reliable connection allows us to use a simple handshaking protocol to ensure state consistency. Event Parser The kernel of the session player is an Event Parser which breaks down session logs into a sequence of event records and retrieves the type and parameter of each event. There are a finite number of event types. As shown in Section 4.1, for each type, the format of the events is a simple regular language. This makes the manual coding of the Event Parser that we decided to use much simpler than using a lexical analyzer generator such as lex. 5.3 Session Player Our multithreaded Session Player performs well; however, careful attention must be paid to the "atomicity" of replaying remote events. Aborting an event in the middle will result in an inconsistent state. All events we are interested in (i.e., those that we capture and relay to the ADL client that's in the slave Mode) finish in a finite and brief amount of time, so we simply use a protocol that ensures the completion of the current event. 5.4 User Interface (Receiver Event Notification) Not all user interface events are obvious if not generated by the user. We found that in replaying events received from the remote Master client, major changes were obvious (such things as zooming in or out in the Map Browser window), but minor changes resulting from selecting a button were easily missed by the user. Ideally, every user interface event could be duplicated, down to tracking the movement of the mouse. This way it is easier to follow subtle changes in the user interface state. It's possible to simulate user interface events at this level of granularity, but it requires a formidable amount of user interface resources. In addition, the time required for implementing such a scheme was not available to us in developing a prototype. We instead added a fairly simple layer between the User Event Handler and the Event Parser that, depending on the user interface component in question, gives the user obvious visual clues. For example, when changing mode from Select to Erase in the Map Browser window, the Select button changes its high-lighting to the default color for non-selected components, and the Erase button flashes red and white several times, finally settling on the default color for selected components. This light-weight and modular approach makes it easy to replace or extend the event notification system. Synchronized Web Browser As mentioned in Section 2, the current ADL client also uses a World Wide Web browser to display images (the results of Figure 3: The relationship between the extension modules and the original system queries). In a separate project, we implemented some primitive synchronization between two World Wide Web browsers using JavaScript and Java; the functionality is similar to but much simpler than [WR94, FLF94]. 6 CONCLUSIONS AND FUTURE WORK Our experience in the design and development of AlexHelp! shows that adding collaborative functionality to an existing system as an independent module can be a viable and fast approach. Our approach can be summarized in the following steps: 1. Examine the existing system and based on the analysis, carefully design the extension system. The design work includes but is not limited to the extension architecture and function specification. 2. Identify one or several likely hook points between the original system and the extension. In our case, one of the hook points was the session logging; we augmented the logging facilities to enable distribution of user interface events. 3. Perform detailed design work, including communication interface, state transition protocols and user events re- production, etc. 4. Implement and link the new module to the existing system As mentioned previously, what we have done is a throw-away prototype; some promising features have been left out due to time constraints, which include: The design of a good help wizard requires much analysis of the problems that users may encounter. It might be interesting to add some logic to the help wizard that allowed it to learn from users' requests. ffl Call Forwarding Call forwarding is not a trivial feature. Since TCP/IP sockets don't allow the migration of connection information from system to system, transparent call forwarding (meaning the user should not be aware that the connection has been broken and re-established) involves another layer upon the TCP/IP protocol. ffl Wait Time Estimation An accurate estimation of the expected waiting time is necessary for a practical system. However, finding a suitable model for such estimation could itself be an interesting research area. One such model would involve using priority-based scheduling for the incoming help call requests. Many factors must be considered when determining priority: user class ("premium user", "common user", etc.) accumulated wait time, the difficulty or urgency of the problem (if reasonable measures can be es- tablished), etc. Scheduling in this fashion is similar to process scheduling in an operating system. Multicast Support Multicast support will become more important as the user group grows larger. The incorporation of multi-cast capability involves modification of the collaboration model and mode control protocol. ffl Client Side Action Analysis Users' session logs reflect the actions they've performed. This data is valuable for the analysis of how the digital library is being used. Data mining may be added to ease the task of analyzing such a potentially large volume of data. Currently ADL is only accessed by a small testing com- however, security issues will become important when the service is open to public. ffl Multimedia Support Multimedia support is not difficult to plug in as several separate channels within the Communication Layer. However, since users can access the library via different links, ranging from high speed network connections to low speed modem connections, users should be able to control the bandwidth by switching on only the multi-media features their connection can accommodate. ACKNOWLEDGMENTS The authors thank Vinod Anupam for his stimulating discussions which lead to the idea of online help studied in this paper; Linda Hill, Mary Larsgaard, and the ADL implementation team, in particular, Nathan Freitas and Kevin Lovette, for their comments and help in the specification and implementation of AlexHelp!; Linda Hill and Mary-Anna Rae for their comments on an earlier version of this paper. --R Automated assistance for the telemeeting lifecycle. Alexandria digital library: Rapid prototype and metadata schema. Extending www for synchronous collaboration. Real time groupware as a distributed system: Concurrency control and its effect on the interface. Eight challenges for developers. JDK 1.1. Java Shared Data Toolkit The impact of CSCW on database technology. A synchronous collaboration tool for World-Wide Web --TR Groupware and social dynamics Computer-Supported Cooperative Work Real time groupware as a distributed system Automated assistance for the telemeeting lifecycle Java object-sharing in Habanero Alexandria Digital Library --CTR Marcos Andr Gonalves , Edward A. Fox , Layne T. Watson , Neill A. Kipp, Streams, structures, spaces, scenarios, societies (5s): A formal model for digital libraries, ACM Transactions on Information Systems (TOIS), v.22 n.2, p.270-312, April 2004

online support;digital library;user interface;online help desk;collaboration

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Time and Cost Trade-Offs in Gossiping.

Each of n processors has a value which should be transmitted to all other processors. This fundamental communication task is called gossiping. In a unit of time every processor can communicate with at most one other processor and during such a transmission each member of a communicating pair learns all values currently known to the other.Two important criteria of efficiency of a gossiping algorithm are its running time and the total number of transmissions. Another measure of quality of a gossiping algorithm is the total number of links used for transmissions. This is the minimum cost of a network which can support the gossiping algorithm. We establish trade-offs between the time T of gossiping and the number C of transmissions and between the time of gossiping and the number L of links used by the algorithm. For a given T we construct gossiping algorithms working in time T, with parameters C and L close to optimal.

Introduction Gossiping (also called all-to-all broadcasting) is one of the fundamental tasks in network communication. Every node of a network (processor) has a piece of information (value) which has to be transmitted to all other nodes, by exchanging messages along the links of the network. Gossiping algorithms have been extensively studied, especially in the last twenty years; see the comprehensive surveys [5, 8] of the domain. The classical communication model, already used in the early papers on gossiping [1, 2, 3, 7, 14], is called the 1-port full-duplex model. Communication is synchronous. In a single round (lasting one unit of time) every node can communicate with at most one neighbor and during such a transmission communicating nodes exchange all values they currently know. Two important criteria of efficiency of a gossiping algorithm are its running time (the number of communication rounds) and the total number of transmissions (calls). The latter is a measure of cost of the algorithm, assuming unit charge per call. The minimum time of gossiping in a complete n-node network was the first problem in this domain, studied in the fifties [2, 14]. It was proved to be dlog ne for even n and dlog ne On the other hand, the minimum number of calls in gossiping is 2n \Gamma 4, for any n ? 3 (cf. [1, 7]). Another measure of quality of a gossiping algorithm is the total number of links used for communication. This is the minimum cost of a network which can support the algorithm, measured by the number of links in the network. This can be also viewed as a measure of the cost of implementing the algorithm, a fixed cost associated with network design, rather than the cost associated with each run. Clearly, the sparsest network supporting gossiping is a tree and thus the minimum number of links is n \Gamma 1. It turns out that the above criteria of efficiency are incompatible: it is impossible to minimize time and the number of calls or to minimize time and the number of links used by the algorithm, simultaneously. If every gossiping algorithm working in time r must have both the number of calls and the number of links used for communication equal to r2 as every node has to communicate in every round with a different node, in order to double its knowledge. On the other hand, Labahn [11] proved that the minimum running time of a gossiping algorithm with the number of calls 2n \Gamma 4 is 2dlog ne \Gamma 3, almost a double of the absolute minimum time. (An earlier proof of this fact, published in [15], was incorrect.) Likewise, in order to minimize the number of links used for communication, we must allow larger gossiping time. Labahn [10] proved that the minimum gossiping time in a tree is at least 2dlog ne \Gamma 1, again almost a double of the absolute minimum time. These results indicate the existence of time vs. cost trade-offs in gossiping, where cost is measured either by the number C of calls or by the number L of links used for communication. Establishing these trade-offs is the main goal of the present paper. For a given T (ranging from log n to 2log n) we show upper and lower bounds on the minimum cost of gossiping in time T . The algorithms yielding our upper bounds are generalizations of known gossiping schemes that minimized separately either the running time or the cost (cf. [1, 2, 3, 7, 11, 14, 15]). While these classical algorithms were either fast but costly or cheap but slow, it turns out that they can be combined to yield almost optimal cost for any given running time. However, the main contribution of this paper are lower bounds on the minimum cost of gossiping for a given running time, that closely match the performance of our respective algorithms. This is for the first time that the full spectrum of relations between the time and the cost of gossiping is investigated. Each of our bounds is useful for a different range of values of the running time and cost. If the running time is ne + t(n), we show an upper bound 2n + O( nlog n ) on the number of calls, which closely matches the lower ) following from [12] . These bounds are useful for small t(n), i.e., when the running time is small. If the running time is ne \Gamma r(n), we show an upper bound 2n + O(r(n)2 r(n) ) and a lower bound log These bounds are useful for small r(n), i.e., when the running time is larger. Here are a few consequences of the above results. Let the running time T of gossiping be equal to dlog ne+ t(n). Let C denote the minimum number of calls in time T . The following sequence of bounds shows how C gradually decreases from \Theta(nlog n) to the asymptotically optimal range 2n + o(n), as restrictions on T are being relaxed. If t(n) is constant then C 2 \Theta(nlog n). If If t(n) - log log n \Gamma d, for a constant d, then C 2 O(n). If For medium range values of the running time T we obtain the following bounds on the minimum number of calls: If log Finally, if we want to keep the number of calls very small, time has to increase significantly: If We also establish trade-offs between the time T of gossiping and the minimum number L of links used for communication. For medium and large values of T the optimum values of L are roughly one-half of the values of C for the same time. In this range we get bounds that are even tighter than in the case of the number C of calls. For example, if log n) and L log n ). For small values of we obtain the upper bound n + O( nlog n on L but our lower bound leaves a larger gap than before: we show that if t(n) - c log log log n, for It remains open, for example, if L 2 \Omega\Gamma nlog n) for constant t(n). The latter bound should be contrasted with a result of Grigni and Peleg [6], concerning broadcasting. They showed that the minimum number of links in an n-node network supporting broadcasting from any node in a given time T is extremely sensitive to the value of is a power of 2, broadcasting in time log n links, while broadcasting in time log can be performed in a network with O(n) links. Our bound shows that this is not the case for gossiping: in particular, gossiping in time log const cannot be performed in a network with a linear number of links. It turns out that the problem of minimizing the cost of gossiping with a given running time has a different flavor in the case of the number of calls and of the number of links. While the same algorithms provide upper bounds in both cases, the techniques used to prove lower bounds are different and results concerning one of these performance measures do not seem to imply meaningful bounds for the other, in any straightforward way. The paper is organized as follows. In section 2 we introduce the terminology and state some preliminary results used in the sequel. Section 3 is devoted to the description of a class of gossiping algorithms and computing their running time, number of calls and number of links used for communication. These results yield upper bounds on the minimum cost of gossiping with a given running time. In section 4 we establish lower bounds on the number of calls in gossiping with a given running time. In section 5 we give lower bounds on the number of links used in gossiping with a given running time. In section 6 we derive consequences of previous results by applying them with appropriate parameter values. Finally, section 6 contains conclusions and open problems. preliminaries The set of communicating nodes is denoted by X and its size is denoted by n. A calling scheme S on the set X is a multigraph on X whose edges are labeled with natural numbers t, so that edges sharing a common node have different labels. Edges with label i represent calls made in the ith time unit. The number of labels is called the running time of the scheme and the number of edges is called the number of calls of the scheme. The corresponding multigraph is called the graph of calls of the scheme S. The underlying graph of a calling scheme S is the simple graph on the set X of nodes in which adjacent nodes are those joined by at least one edge in S. This is the minimal network that supports the scheme S. The number of edges in the underlying graph of S is called the number of links used by S. Upon completion of S the node v knows the value of the node w if there exists an ascending path from w to v in S, i.e., a path with increasing labels on edges. The set of nodes who know the value of v upon completion of S is denoted by K(v) and the set of nodes whose value v knows upon completion of S is denoted by K \Gamma (v). If the calling scheme S is called a gossiping scheme or gossiping algorithm. The (total) knowledge upon completion of the calling scheme S is the number (v)j. The knowledge after i rounds is at most for every node v. The knowledge at the end of a gossiping scheme is n 2 . Lemma 2.1 1. If the calling scheme has k calls then jK(v)j for every node v. 2. If the running time of a calling scheme is t then jK(v)j node v. Lemma 2.2 If then the time required for the remaining nodes to learn the value of v is at least log n \Gamma log k. Proof: One of the k informed nodes has to inform at least n\Gammak other nodes which requires time at least log n All logarithms are with base 2. The notation O,\Omega and \Theta is standard. We use o(f(n)) (resp. !(f(n))) to denote the class of functions g(n), such that g(n) converges to 0, as grows. 3 Gossiping algorithms and upper bounds In this section we present a class of gossiping algorithms that provide good time and cost trade-offs, both in the case when cost is measured by the number of calls and when it is measured by the number of links used for communication. Two important graphs will be used in the construction of our schemes. The first is the k-dimensional hypercube H k . This is the graph on 2 k nodes labeled with all binary sequences of length k. Nodes are adjacent iff their labels differ in exactly one position. Nodes whose labels differ in the jth position are called j-neighbors. The second graph is the k-broadcasting tree B k . It is defined by induction on k. B 0 is a single node v. B k+1 is obtained from B k by attaching a different new node to every node of . The set of all new edges is called the 1)th layer in B k+1 . The initial node v is called the root of the broadcasting tree. Hypercubes and broadcasting trees are important for gossiping. Giving label j to edges of the hypercube H k joining j-neighbors yields a gossiping scheme with the smallest running time k. The cost of this scheme, however, is very large: it uses k2 k\Gamma1 calls and k2 links. On the other hand, broadcasting trees yield gossiping schemes with small cost but large time. Replace every edge in layers by two edges: one with label the other with label k 1). Give label k to the edge in layer 1. The obtained gossiping scheme first gathers all values in the root and then broadcasts them all to all nodes. Its running time is its cost is very low: if is the number of nodes, it has the optimal number links and it uses 2n \Gamma 3 calls, only one call more than the absolute minimum. In order to save gossiping time at a given cost or to lower cost with a given running time, it is advantageous to use a combination of the two above schemes. Let ne \Gamma 1. Let r - k and We describe the gossiping algorithm COT(n; r). (COT stands for Cube of Trees.) Consider the hypercube H r and let each of its nodes be the root of a broadcasting tree B s . Trees rooted at distinct nodes of H r are disjoint. There are 2 k nodes in all trees. Attach each of the remaining x nodes to a distinct node in one of the trees. Define the set of edges incident to these nodes to be the Replace each edge of layers by two edges with labels s Finally give label s to edges of the hypercube H r joining i-neighbors. The above described gossiping scheme works as follows: first information from all nodes of the tree rooted at a given node of the hypercube is gathered in this node. Then gossiping is executed inside the hypercube H r among all its nodes. At this point all nodes of the hypercube know all values. Finally each node of the hypercube broadcasts the complete information to all nodes of the tree rooted at it. The underlying graph of the scheme COT(n; r) is the undirected version of the graph H r;s used in [6] for broadcasting. Theorem 3.1 The gossiping algorithm COT(n; r) has running time uses links. Proof: Gathering information in nodes of H r takes time gossiping in H r takes time r and broadcasting complete information in trees takes time total of Gathering information in nodes of H r uses 2 r gossiping in H r uses r2 calls and broadcasting complete information in trees uses again 2 r calls, for a total of The number of links in the hypercube H r is r2 r\Gamma1 , the total number of links in all trees is a total of The above theorem yields upper bounds on the cost of gossiping with a given running time. It will be convenient for our purposes to formulate them in two versions: Corollary 3.1 For any functions t; r : N ! N , such that t(n); r(n) - dlog ne, there exists a gossiping algorithm 1. with running time ne \Gamma r(n), number of calls C 2 2n +O(r(n)2 r(n) ) and using 2. with running time ne number of calls C using links. Proof: 1. Straightforward. 2. Use part 1. for The above corollary shows that there exists a gossiping algorithm whose time and cost are both asymptotically optimal, i.e., whose running time is log n) and which uses links. To this end it suffices to take, e.g., However, the results of the following sections will enable us to establish time and cost trade-offs more precisely. Lower bounds on the number of calls In this section we give two lower bounds on the number of calls in gossiping with a given running time. Each of them provides meaningful consequences for a different range of time and cost values. The first bound follows directly from a result of Labahn [12] and is useful for small values of the running time. Theorem 4.1 Every gossiping algorithm with running time calls. The next theorem yields lower bounds on the number of calls in gossiping that are useful when the running time is larger. We first prove two lemmas. Lemma 4.1 If a calling scheme has the running time at most t and its graph of calls is a tree then: 1. there exists a node v such that jK(v)j - t 2. there exists a node v such that Proof: We prove only the first part of the lemma: the second part is analogous. Call a node terminal if there is no ascending path of length larger than 1, starting from v. It suffices to prove that there exists a terminal node v. Indeed, for such a node, K(v) consists of v itself and of its neighbors in the tree of calls. The desired inequality follows from the fact that the number of neighbors of a node in the graph of calls cannot exceed the running time of the calling scheme. Choose any node w 0 and suppose that it is not terminal. Choose any ascending path (w 2. If w 2 is terminal, we are done, if not, choose any ascending path (w length 2, and so on. Since labels in each path are strictly increasing and the graph of calls is a tree, at every step at least one new node is visited. Thus the process must terminate at some node w k which has to be terminal. 2 Lemma 4.2 If a calling scheme on n nodes has the running time at most t and uses at most 1. there exists a node v such that jK(v)j - t 2. there exists a node v such that Proof: Again we prove only the first part of the lemma. Suppose that S is a calling scheme satisfying the assumptions but violating assertion 1. Let a 1 ; :::; a k be the numbers of nodes in components of the graph of calls of S. No component has a node v such that jK(v)j - t + 1. It follows from lemma 4.1 that none of the components can be a tree hence the ith component must have at least a i edges. Hence the total number of edges in the graph of calls is at least n, contradicting the assumption on the number of calls in S. 2 Theorem 4.2 Every gossiping algorithm with running time log calls. Proof: Let t be the largest integer such that less than n calls are placed before round t. Let S be the calling scheme consisting of all calls of S with labels at most t \Gamma 1. Lemma 4.2 implies that after time there is a node v such that jK(v)j - 2log n. (Here K(v) is taken with respect to the calling scheme S .) By lemma 2.2 the additional time required for all nodes to learn the value of v is at least log n Hence and consequently be the calling scheme consisting of all calls of S with labels at most t. (The number of calls in S 1 is at least n.) Lemma 2.1 implies that after the first t rounds, for every node v 2 X. (Here sets K \Gamma (v) are taken with respect to the calling scheme S 1 .) consider the calling scheme S 2 consisting of the first a(n) calls placed after round t (order calls in the same round arbitrarily). Lemma 2.1 implies that, for every node v 2 X, taken with respect to S 2 . Thus, upon completion of all calls in schemes S 1 and S 2 , for every node v 2 X. Now at most remain to be placed. Denote by S 3 the scheme consisting of these remaining calls. By lemma 4.2 there exists a node w, such that now taken with respect to the scheme S 3 . It follows that upon completion of all calls in schemes S 1 , S 2 and S 3 , i.e., at the end of the scheme S, node w knows the values of at most nodes. Since S is a gossiping scheme, we must have whence log This concludes the proof. 2 5 Lower bounds on the number of links In this section we establish two lower bounds on the number of links used by a gossiping scheme with a given running time. The first bound concerns the case when the running time is small. Theorem 5.1 Every gossiping algorithm with running time T - log n+c log log log n, where uses L 2 !(n(log log n) d ) links, where d Before proving the theorem we fix some additional terminology and prove several technical lemmas. Consider a calling scheme with running time T . Let log log log n, c ! 1. Suppose that the number of links used by this scheme is log log log n, for sufficiently large n. We will prove that the considered calling scheme is not a gossiping scheme. Suppose it is. A node v is called after round i of the scheme if jK \Gamma (v)j is at most 1 a node that is not weak is called strong. A call between nodes v and w in round i is said to be ff-increasing round i is at most ff times larger than this round. Let In every round i consider the following classes of calls: A - Calls between weak nodes, ffl)-increasing calls not belonging to the class A, remaining calls. The idea of the proof is to show that in many rounds there are few nodes that are either or participate in calls of class C and consequently the increase of knowledge in these rounds is too slow to enable achieving knowledge n 2 upon completion of the scheme. Among our arguments many hold only for sufficiently large n. This does not cause any problems, since the result is of asymptotic nature. We skip the phrase "for sufficiently large n" for the sake of brevity. We start with a lower bound on the number of strong nodes. Lemma 5.1 In every round there are at least 3 strong nodes. Proof: After every round i the knowledge K is at least because in the remaining log rounds knowledge can increase at most 2 log n+f(n)\Gammai times and the final knowledge must be n 2 . Let p be the number of strong nodes and the number of weak nodes after the ith round. After the ith round the knowledge K is at most p2 i p)2 i\Gammaf (n)\Gamma2 , hence which implies :The aim of the next two lemmas is to give an upper bound on the size of the class C. Define the forbidden distance to be the maximum number k such that if a call of class C has been placed on a link in round i then no call of this class is placed on this link in rounds Lemma 5.2 The forbidden distance is at least 2 f(n)+b(n)+8 . Proof: Suppose that a call of class C has been placed on link )j be the amount of information in each of these nodes after this round. Let l be the minimum positive integer such that a call of class C is placed on link e in round i l. We will show that l ? 2 f(n)+b(n)+8 . Since the call on link e in round i was in the class C, at least one of the nodes v 1 or v 2 was strong after round i \Gamma 1. Thus Consider the increase of the number after round 2. We have the upper bound requiring that v j communicate in every round having maximum and mutually disjoint information. On the other hand, jK w after round inequality was already true after round i. After round l we have hence the increase of the number l is at most In view of inequality 1 the right hand side of the above is at most Since the call in round i+l on link e is in the class C, it is not in the class B and consequently the number must increase in round times. Hence we get which implies and finally l Lemma 5.3 jCj - nlog n Proof: Since the total number of rounds is less than 2log n, there are at most 2log n calls of class C on every link. The total number of links is at most :The next two lemmas show that in many rounds there are many strong nodes that do not participate in calls of class C. Call a round essential if there are at most n calls of class C in this round. Lemma 5.4 At least Trounds are essential. Proof: Otherwise more than Trounds would have more than n calls of class C, for a total of more than2 log n \Delta n which contradicts lemma 5.3. 2 Lemma 5.5 In every essential round there are at least n strong nodes that do not participate in calls of class C. Proof: By lemma 5.1 there are at most n(1 \Gamma 3 nodes in every round. By definition there are at most n calls of class C in every essential round. At most n nodes can participate in these calls. Hence the total number of nodes that are either weak or participate in a call of class C is at most in every essential round. 2 The next two lemmas show that in many rounds the rate of knowledge increase can be bounded strictly below 2. Call a pair of nodes fv; wg red in round i if (w)j is at least 1 round sum increases at most 2 \Gamma ffl times in round i; otherwise call the pair fv; wg white in round i. Lemma 5.6 In every essential round there are at least n red pairs of nodes. Proof: Fix an essential round i. A strong node v that does not participate in a call of class C, either participates in a call of class B or does not communicate at all in round i. By lemma 5.5, there are either at least n nodes of the first type or of the second type. In the first case there are at least n calls in the class B because every such call involves at least one strong node (otherwise it would be in the class A). All pairs of nodes in these calls are red, which proves the lemma in this case. In the second case partition nodes that do not communicate in the ith round into disjoint pairs arbitrarily. does not increase at all in such pairs in the ith round and at least n pairs contain a strong node in this case, hence they are red. 2 Lemma 5.7 In every essential round the total knowledge K increases at most times; where Proof: For simplicity assume that the number of nodes is even, it will be clear how to modify the argument otherwise. Fix an essential round i. By lemma 5.6 there are at least disjoint red pairs in round i. For every such pair fv; wg, after round and the increase of (w)j in this round is at most 2 \Gamma ffl times. For pairs fv; wg that are white in round i, and the increase of (w)j is at most 2 times. We want to establish an upper bound on the rate of knowledge increase in round i. We will compute this rate as a fraction R whose numerator is the sum of disjoint pairs of nodes after round i and the denominator is the corresponding sum before round i. The value of R canot decrease if the number of red pairs is decreased to n and the sum lowered to 1 in every red pair, while the number of white pairs is increased to n\Gamma n and the sum increased to 2 i in every white pair. Also R cannot decrease if we assume that the increase of times in red pairs and 2 times in white pairs. Hence we get Denoting simplifying we get 4x 4x and finally . 2 Proof of theorem 5.1: Denote, as before, . By lemmas 5.4 and 5.7, knowledge increases at most times in at least 1log n rounds. In all remaining rounds it increases at most 2 times. Hence in order to show that our scheme is not a gossiping scheme it suffices to show i.e., log log log n, for log log log n, for d log n: g(n) . The latter inequality implies h(n) - 2f(n). Since e , we have 2:5 for sufficiently large n and thus In view of h(n) - 2f(n) we have which implies inequality 2. 2 The last result of this section gives a meaningful lower bound on the number of links when the running time is in the medium or large range. Theorem 5.2 Every gossiping algorithm with running time T - 2log uses log n links. Proof: We may assume that the conclusion is trivial. Suppose that L - 16log n . Take a spanning tree of the underlying graph, with root k, diameter at most 2log n and maximum degree at most 2log n. Such a tree must exist for the gossiping to be completed in time less than 2log n. Color all links of this tree black and all other links (at most 2 r(n) red links to the tree one by one, each time recoloring red those black links which appear in a newly created cycle. If the link fv; wg is added, this causes recoloring red links on the paths joining v with k and w with k in the tree. (Some of them may have been recolored already previously.) Hence adding a new red link causes recoloring at most 2log n black links. After adding at most 2 r(n) red links, the total number of red links at the end of the recoloring process is at most which is less than 2 r(n)\Gamma2 for sufficiently large n, in view of Since links that are red at the end of the recoloring process are exactly those situated in cycles in the underlying graph, this graph has z ! 2 r(n)\Gamma2 nodes situated in cycles. Hence there exists a tree D attached to only one node d in some cycle, such that z Case 1. 2n nThe value of some node v in D reaches the node d after time larger than log 2n 1. Broadcasting the value of v from d to all nodes outside of D requires time at least log n= log n \Gamma 1. Hence the total time of gossiping exceeds 2log Case 2. jDj ? nSince the maximum degree of D is at most 2log n, the tree D contains a subtree Y , such that 2n . The rest of the argument is as in Case 1, with D replaced by Y . 2 6 Discussion We have two pairs of bounds on the minimum number of calls C in gossiping with a given running time T . If ne log ). The first pair of bounds is useful for small t(n), e.g., when t(n) 2 O(log log n), i.e., when gossiping time is small. They yield the following corollary showing how C gradually decreases from \Theta(nlog n) to the asymptotically optimal range 2n + o(n), as restrictions on T are being relaxed. Corollary 6.1 If ne 1. If t(n) is constant then C 2 \Theta(nlog n). 2. If t(n) 2 log log 3. If t(n) - log log n \Gamma d, for a constant d, then C 2 O(n). 4. If t(n) 2 log log n The lower bound C 2 \Omega\Gamma nlog n ), following from [12], becomes trivial when t(n) ? log log n. For even larger values of gossiping time our second pair of bounds can be applied. For example, it gives a fairly precise estimate of the minimum number of calls when the running time is in the medium range fflog n, 2. Corollary 6.2 If the running time of a gossiping algorithm is then C 2 2n +O(n 2\Gammaff log n) and C log The next corollary corresponds to the situation when the gossiping time is fairly large. In this case it is more natural to reverse the problem: what is the minimum running time of gossiping when the number of calls has to be kept very small? Corollary 6.3 If the number of calls in a gossiping algorithm is is polylogarithmic in n, then its running time T is 2log Proof: Suppose not, and let 2\Omega\Gamma367 n). Then r(n) - dlog n, for some constant d and C log We next turn attention to the trade-off between the time T and the number of links L. For small values of T the gap between our upper and lower bounds is larger than in the previous case. Corollary 3.1 and theorem 5.1 imply, for example, that if a constant, then L 2 O(nlog n) and L 2 !(n(log log n) d ), for d ! 1. It remains open if n) in this case. The last pair of bounds, applicable for larger values of gossiping time follows from corollary 3.1 and theorem 5.2. In this case L 2 log n ). For the medium range of gossiping time fflog n, gives an even more precise estimate of L than previously of C. Corollary 6.4 If the running time of a gossiping algorithm is then log n) and L log n Finally, a result similar to corollary 6.3 holds for the number of links. Corollary 6.5 If the number of links used by a gossiping algorithm is c(n) is polylogarithmic in n, then its running time T is 2log 7 Conclusion We established upper and lower bounds on the minimum number of calls and the minimum number of links used by a gossiping scheme with a given running time. Our algorithms, which turned out to be cost-efficient for the whole range of running time values, follow the same simple pattern: gather information in nodes of a hypercube of appropriately chosen size using a separate broadcasting tree for each node, then gossip in the hypercube in minimal time and finally broadcast complete information to all remaining nodes, using again the same broadcasting trees. The tree part of the scheme uses few calls and few links but a lot of time, as it is executed twice, while the hypercube part is fast but uses many calls and many links. Thus a suitable balance between these parts must be maintained to get low cost for a given running time. Our bounds leave very small gaps. For example, if our upper bound on C is log n) and the lower bound is 2n log leaving a gap within a factor of O(log 3 n) in the part of the number of calls exceeding the absolute minimum 4. In the case of the number of links L, our bounds are even tighter for this range of running time. For the same value our upper bound on L is log n) and the lower bound is n log n ), leaving a gap within a factor of O(log 2 n) in the part of the number of links exceeding the absolute minimum n \Gamma 1. Further tightening of these bounds, for all values of running time, remains a natural open problem yielded by our results. We do not know, for example, if it is possible to gossip in time 3log n using 2n n) calls and/or n n) links. It also remains open what is the minimum value of L when const. We conjecture that L 2 \Theta(nlog n) in this case. Another interesting problem is to evaluate the complexity of finding the exact value of the minimum cost of gossiping with a given running time. Given n and T , can the minimum number of calls C or the minimum number of links L be found in polynomial time? In many papers (cf. [4, 9, 10]) gossiping was studied for specific important networks, such as trees, grids or hypercubes, and the time or the number of calls were minimized separately. It would be interesting to extend our study by investigating time vs. number of calls trade-offs in gossiping for these networks as well. Also communication models other than the classical 1-port full-duplex model (cf., e.g., [9]), could be considered in this context. --R Gossips and telephones Communication patterns in task-oriented groups A problem with telephones Gossiping in grid graphs Methods and problems of communication in usual net- works Tight bounds on minimum broadcast networks A cure for the telephone disease A survey of gossiping and broadcasting in communication networks Fast gossiping for the hypercube The telephone problem for trees Kernels of minimum size gossip schemes Some minimum gossip graphs The distribution of completion times for random communication in a task-oriented group Time and call limited telephone problem --TR --CTR Francis C.M. Lau , Shi-Heng Zhang, Fast Gossiping in Square Meshes/Tori with Bounded-Size Packets, IEEE Transactions on Parallel and Distributed Systems, v.13 n.4, p.349-358, April 2002 Francis C. M. Lau , S. H. Zhang, Optimal gossiping in paths and cycles, Journal of Discrete Algorithms, v.1 n.5-6, p.461-475, October

gossiping;lower bounds;algorithm

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High-level design verification of microprocessors via error modeling.

A design verification methodology for microprocessor hardware based on modeling design errors and generating simulation vectors for the modeled errors via physical fault testing techniques is presented. We have systematically collected design error data from a number of microprocessor design projects. The error data is used to derive error models suitable for design verification testing. A class of basic error models is identified and shown to yield tests that provide good coverage of common error types. To improve coverage for more complex errors, a new class of conditional error models is introduced. An experiment to evaluate the effectiveness of our methodology is presented. Single actual design errors are injected into a correct design, and it is determined if the methodology will generate a test that detects the actual errors. The experiment has been conducted for two microprocessor designs and the results indicate that very high coverage of actual design errors can be obtained with test sets that are complete for a small number of synthetic error models.

INTRODUCTION It is well known that about a third of the cost of developing a new microprocessor is devoted to hardware debugging and testing [25]. The inadequacy of existing hardware verification methods is graphically illustrated by the Pentium's FDIV error, which cost its manufacturer an estimated $500 million. The development of practical verification methodologies for hardware verification has long been handicapped by two related problems: (1) the A preliminary version of this paper was presented in [4] at the 1997 IEEE International High Level Design Validation and Test Workshop, Oakland, California, November 14-15, 1997. Permission to make digital or hard copies of part or all of this work for personal or classroom use is granted without fee provided that copies are not made or distributed for profit or direct commercial advantage and that copies show this notice on the first page or initial screen of a display along with the full citation. Copyrights for components of this work owned by others than ACM must be hon- ored. Abstracting with credit is permitted. To copy otherwise, to republish, to post on servers, to redistribute to lists, or to use any component of this work in other works, requires prior specific permission and/or a fee. Permissions may be requested from Publications Dept., ACM Inc., 1515 Broadway, New York, NY 10036 USA, fax +1 (212) 869-0481, or [email protected]. - 1998 by the Association for Computing Machinery, Inc. Al-Asaad, J. P. Hayes, T. Mudge, and R. B. Brown lack of published data on the nature, frequency, and severity of the design errors occurring in large-scale design projects; and (2) the absence of a verification methodology whose effectiveness can readily be quantified. There are two broad approaches to hardware design verification: formal and simula- tion-based. Formal methods try to verify the correctness of a system by using mathematical proofs [32]. Such methods implicitly consider all possible behavior of the models representing the system and its specification, whereas simulation-based methods can only consider a limited range of behaviors. The accuracy and completeness of the system and specification models is a fundamental limitation for any formal method. Simulation-based design verification tries to uncover design errors by detecting a cir- cuit's faulty behavior when deterministic or pseudo-random tests (simulation vectors) are applied. Microprocessors are usually verified by simulation-based methods, but require an extremely large number of simulation vectors whose coverage is often uncertain. Hand-written test cases form the first line of defense against bugs, focusing on basic functionality and important corner (exceptional) cases. These tests are very effective in the beginning of the debug phase, but lose their usefulness later. Recently, tools have been developed to assist in the generation of focused tests [13,20]. Although these tools can significantly increase design productivity, they are far from being fully automated. The most widely used method to generate verification tests automatically is random test generation. It provides a cheap way to take advantage of the billion-cycles-a-day simulation capacity of networked workstations available in many big design organizations. Sophisticated systems have been developed that are biased towards corner cases, thus improving the quality of the tests significantly [2]. Advances in simulator and emulator technology have enabled the use of very large sets as test stimuli such as existing application and system software. Successfully booting the operating system has become a basic quality requirement [17,25]. Common to all the test generation techniques mentioned above is that they are not targeted at specific design errors. This poses the problem of quantifying the effectiveness of a test set, such as the number of errors covered. Various coverage metrics have been proposed to address this problem. These include code coverage metrics from software testing [2,7,11], finite state machine coverage [20,22,28], architectural event coverage [22], and observability-based metrics [16]. A shortcoming of all these metrics is that the relationship between the metric and the detection of classes of design errors is not well understood A different approach is to use synthetic design error models to guide test generation. This exploits the similarity between hardware design verification and physical fault test- ing, as illustrated by Figure 1. For example, Al-Asaad and Hayes [3] define a class of design error models for gate-level combinational circuits. They describe how each of these errors can be mapped onto single-stuck line (SSL) faults that can be targeted with standard automated test pattern generation (ATPG) tools. This provides a method to generate tests with a provably high coverage for certain classes of modeled errors. A second method in this class stems from the area of software testing. Mutation testing [15] considers programs, termed mutants, that differ from the program under test by a single small error, such as changing the operator from add to subtract. The rationale for the approach is supported by two hypotheses: 1) programmers write programs that are close to High-Level Design Verification of Microprocessors via Error Modeling - 3 correct ones, and 2) a test set that distinguishes a program from all its mutants is also sensitive to more complex errors. Although considered too costly for wide-scale industrial use, mutation testing is one of the few approaches that has yielded an automatic test generation system for software testing, as well as a quantitative measure of error coverage (mutation score) [24]. Recently, Al Hayek and Robach [5] have successfully applied mutation testing to hardware design verification in the case of small VHDL modules. This paper addresses design verification via error modeling and test generation for complex high-level designs such as microprocessors. A block diagram summarizing our methodology is shown in Figure 2. An implementation to be verified and its specification are given. For microprocessors, the specification is typically the instruction set architecture (ISA), and the implementation is a description of the new design in a hardware description language (HDL) such as VHDL or Verilog. In this approach, synthetic error models are used to guide test generation. The tests are applied to simulated models of both the implementation and the specification. A discrepancy between the two simulation outcomes indicates an error, either in the implementation or in the specification. Section 2 describes our method for design error collection and presents some preliminary design error statistics that we have collected. Section 3 discusses design error modeling and illustrates test generation with these models. An experimental evaluation of our methodology and of the error models is presented in Section 4. Section 5 discusses the results and gives some concluding remarks. 2. DESIGN ERROR COLLECTION Hardware design verification and physical fault testing are closely related at the conceptual level [3]. The basic task of physical fault testing (hardware design verification) is to generate tests that distinguish the correct circuit from faulty (erroneous) ones. The class of faulty Prototype system Operational system Design Manufacturing Verification tests Physical fault tests Design errors Physical faults Design development Field deployment model Fault model 1. Correspondence between design verification and physical fault testing. residual design errors D. Van Campenhout, H. Al-Asaad, J. P. Hayes, T. Mudge, and R. B. Brown circuits to be considered is defined by a logical fault model. Logical fault models represent the effect of physical faults on the behavior of the system, and free us from having to deal with the plethora of physical fault types directly. The most widely used logical fault model, the SSL model, combines simplicity with the fact that it forces each line in the circuit to be exercised. Typical hardware design methodologies employ hardware description languages as their input medium and use previously designed high-level modules. To capture the richness of this design environment, the SSL model needs to be supplemented with additional error models. The lack of published data on the nature, frequency, and severity of the design errors occurring in large-scale projects is a serious obstacle to the development of error models for hardware design verification. Although bug reports are collected and analyzed internally in industrial design projects the results are rarely published. Examples of user-oriented bug lists can be found in [21,26]. Some insight into what can go wrong in a large processor design project is provided in [14]. The above considerations have led us to implement a systematic method for collecting design errors. Our method uses the CVS revision management tool [12] and targets ongoing design projects at the University of Michigan, including the PUMA high-performance microprocessor project [9] and various class projects in computer architecture and VLSI 2. Deployment of proposed design verification methodology. Design error models Test generator Implementation simulator Specification simulator Equal? Diagnose Specification Unverified design Verified design CVS revision database Unknown actual error Assisted verification Assisted verification High-Level Design Verification of Microprocessors via Error Modeling - 5 design, all of which employ Verilog as the hardware description medium. Designers are asked to archive a new revision via CVS whenever a design error is corrected or whenever the design process is interrupted, making it possible to isolate single design errors. We have augmented CVS so that each time a design change is entered, the designer is prompted to fill out a standardized multiple-choice questionnaire, which attempts to gather four key pieces of information: (1) the motivation for revising the design; (2) the method by which a bug was detected; (3) a generic design-error class to which the bug belongs, and (4) a short narrative description of the bug. A uniform reporting method such as this greatly simplifies the analysis of the errors. A sample error report using our standard questionnaire is shown in Figure 3. The error classification shown in the report form is the result of the analysis of error data from several earlier design projects. Design error data has been collected so far from four VLSI design class projects that involve implementing the DLX microprocessor [19], from the implementation of the LC-2 microprocessor [29] which is described later, and from preliminary designs of PUMA's fixed-point and floating-point units [9]. The distributions found for the various representative design errors are summarized in Table 1. Error types that occurred with very low frequency are combined in the "others" category in the table. (replace the _ with X where MOTIVATION: correction _ design modification _ design continuation _ performance optimization _ synthesis simplification _ documentation BUG DETECTED BY: _ inspection _ compilation simulation _ synthesis Please try to identify the primary source of the error. If in doubt, check all categories that apply. _ verilog syntax error _ conceptual error combinational logic: wrong signal source _ missing input(s) _ unconnected (floating) input(s) _ unconnected (floating) _ conflicting outputs _ wrong gate/module type _ missing instance of gate/module _ sequential logic: _ extra latch/flipflop _ missing latch/flipflop _ extra state _ missing state _ wrong next state _ other finite state machine error _ statement: _ if statement _ case statement _ always statement _ declaration _ port list of module declaration _ expression (RHS of assignment): _ missing term/factor _ extra term/factor _ missing inversion _ extra inversion _ wrong operator _ wrong constant _ completely wrong _ buses: _ wrong bus width _ wrong bit order _ new category (describe below) Used wrong field from instruction 3. Sample error report. 6 - D. Van Campenhout, H. Al-Asaad, J. P. Hayes, T. Mudge, and R. B. Brown 3. ERROR MODELING Standard simulation and logic synthesis tools have the side effect of detecting some design error categories of Table 1, and hence there is no need to develop models for those particular errors. For example a simulator such as Verilog-XL [10] flags all Verilog syntax errors (category 9), declaration statement errors (category 12), and incorrect port lists of modules (category 16). Also, logic synthesis tools, such as those of Synopsys, usually flag all wrong bus width errors (category 10) and sensitivity-list errors in the always statement (category 13). To be useful for design verification, error models should satisfy three requirements: (1) tests (simulation vectors) that provide complete coverage of the modeled errors should also provide very high coverage of actual design errors; (2) the modeled errors should be amenable to automated test generation; (3) the number of modeled errors should be relatively small. In practice, the third requirement means that error models that define a number of error instances linear, or at most quadratic in the size of the circuit are preferred. The error models need not mimic actual design bugs precisely, but the tests derived from complete coverage of modeled errors should provide very good coverage of actual design bugs. 3.1 Basic error models A set of error models that satisfy the requirements for the restricted case of gate-level logic circuits was developed in [3]. Several of these models appear useful for the higher-level (RTL) designs found in Verilog descriptions as well. From the actual error data in Table 1, we derive the following set of five basic error models: 1. Actual error distributions from three groups of design projects. Design error category Relative frequency [%] 1. Wrong signal source 29.9 28.4 25.0 2. Conceptual error 39.0 19.1 0.0 3. Case statement 4. Gate or module input 11.2 9.8 0.0 5. Wrong gate/module type 12.1 6. Wrong constant 0.4 5.7 10.0 7. Logical expression wrong 8. Missing input(s) 0.0 5.2 0.0 9. Verilog syntax error 10. Bit width error 0.0 2.2 15.0 11. If statement 1.1 1.6 5.0 12. Declaration statement 13. Always statement 0.4 1.4 5.0 14. FSM error 3.1 15. Wrong operator 1.7 0.3 0.0 16. Others 1.1 5.8 25.0 High-Level Design Verification of Microprocessors via Error Modeling - 7 . Bus SSL error (SSL): A bus of one or more lines is (totally) stuck-at-0 or stuck-at- 1 if all lines in the bus are stuck at logic level 0 or 1. This generalization of the standard SSL model was introduced in [6] in the context of physical fault testing. Many of the design errors listed in Table 1 can be modeled as SSL errors (categories 4 and 6). . Module substitution error (MSE): This refers to mistakenly replacing a module by another module with the same number of inputs and outputs (category 5). This class includes word gate substitution errors and extra/missing inversion errors. . Bus order error (BOE): This refers to incorrectly ordering the bits in a bus (category 16). Bus flipping appears to be the most common form of BOE. . Bus source error (BSE): This error corresponds to connecting a module input to a wrong source (category 1). . Bus driver error (BDE): This refers to mistakenly driving a bus with two sources (category 16). Direct generation of tests for the basic error models is difficult, and is not supported by currently available CAD tools. While the errors can be easily activated, propagation of their effects can be difficult, especially when modules or behavioral constructs do not have transparent operating modes. In the following we demonstrate manual test generation for various basic error models. 3.2 Test generation examples Because of their relative simplicity, the foregoing error models allow tests to be generated and error coverage evaluated for RTL circuits of moderate size. We analyzed the test requirements of two representative combinational circuits: a carry-lookahead adder and an ALU. Since suitable RTL tools are not available, test generation was done manually, but in a systematic manner that could readily be automated. Three basic error models are consid- ered: BOEs, MSEs, and BSEs. Test generation for SSLs is discussed in [1,6] and no tests are needed for BDEs, since the circuits under consideration do not have tristate buses. Example 1: The 74283 adder An RTL model [18] of the 74283 4-bit fast adder [30] appears in Figure 4. It consists of a carry-lookahead generator (CLG) and a few word gates. We show how to generate tests for some design error models in the adder and then we discuss the overall coverage of the targeted error models. BOE on A bus: A possible bus value that activates the error is A an unknown value. The erroneous value of A is thus A Hence, we can represent the error by represents the error signal which is 1 (0) in the good circuit and 0 (1) in the erroneous circuit. One way to propagate this error through the AND gate G 1 is to set Hence, we get G and G Now for the module CLG we have X. The resulting outputs are This implies that hence the error is not detected at the primary outputs. We need to assign more input values to propagate the error. If we set C DXXD DXXD DXXD 8 - D. Van Campenhout, H. Al-Asaad, J. P. Hayes, T. Mudge, and R. B. Brown Hence, the error is propagated to S and the complete test vector is (A, B, C (0XX11XX10). On generating tests for all BSEs in the adder we find that just 2 tests detect all 33 detectable BSEs, and a single BSE is redundant as shown above. We further targeted all MSEs in the adder and we found that 3 tests detect all 27 detectable MSEs and proved that a single MSE (G 3 /XNOR) is redundant. Finally, we found that all BOEs are detected by the tests generated for BSEs and MSEs. Therefore, complete coverage of BOEs, BSEs, and MSEs is achieved with only 5 tests. Example 2: The c880 ALU In this example, we try to generate tests for some modeled design errors in the c880 ALU, a member of the ISCAS-85 benchmark suite [8]. A high-level model based on a Verilog description of the ALU [23] is shown in Figure 5; it is composed of six modules: an adder, two multiplexers, a parity unit, and two control units. The circuit has 60 inputs and 26 out- puts. The gate-level implementation of the ALU has 383 gates. The design error models to be considered in the c880 are again BOEs, BSEs, and MSEs (inversion errors on 1-bit signals). We next generate tests for these error models. BOEs: In general, we attempt to determine a minimum set of assignments needed to detect each error. Some BOEs are redundant such as the BOE on B (PARITY), but most BOEs are easily detectable. Consider, for example, the BOE on D. One possible way to activate the error is to set To propagate the error to a primary output, the path through IN-MUX and then OUT-MUX is selected. The signal values needed to activate this path are: Solving the gate-level logic equations for G and C we get: All signals not mentioned in the above test have don't care values. We found that just 10 tests detect all 22 detectable BOEs in the c880 and serve to prove that another 2 BOEs are redundant. 4. High-level model of the 74283 carry-lookahead adder [18]. A G High-Level Design Verification of Microprocessors via Error Modeling - 9 MSEs: Tests for BOEs detect most, but not all, inversion errors on multibit buses. In the process of test generation for the c880 ALU, we noticed a case where a test for an inversion error on bus A can be found even though the BOE on A is redundant. This is the case when an n-bit bus (n odd) is fed into a parity function. Testing for inversion errors on 1-bit signals needs to be considered explicitly, since a BOE on a 1-bit bus is not possible. Most inversion errors on 1-bit signals in the c880 ALU are detected by the tests generated for BOEs and BSEs. This is especially true for the control signals to the multiplexers. 3.3 Conditional error model The preceding examples, as well as prior work on SSL error detection [1,6], show that the basic error models can be used with RTL circuits, and that high, but not complete, error coverage can be achieved with small test sets. These results are further reinforced by our experiments on microprocessor verification (Section which indicate that a large fraction of actual design errors (67% in one case and 75% in the other) is detected by complete test sets for the basic errors. To increase coverage of actual errors to the very high levels needed for design verification, additional error models are required to guide test generation. Many more complex error models can be derived directly from the actual data of Table 1 to supplement the basic error types, the following set being representative: . Bus count error (BCE): This corresponds to defining a module with more or fewer input buses than required (categories 4 and 8). 5. High-level model of the c880 ALU. IN-MUX OUT-MUX A G Cin C Cont ParA ParB F Par-Hi Par-Al Par-Bl Pass-B Usel-G Cout GEN ADDER Pass-H F-add F-and F-xor Al-Asaad, J. P. Hayes, T. Mudge, and R. B. Brown . Module count error (MCE): This corresponds to incorrectly adding or removing a module (category 16), which includes the extra/missing word gate errors and the extra/missing registers. . Label count error (LCE): This error corresponds to incorrectly adding or removing the labels of a case statement (category 3). . Expression structure error (ESE): This includes various deviations from the correct expression (categories 3, 6, 7, 11, 15), such as extra/missing terms, extra/ missing inversions, wrong operator, and wrong constant. . State count error (SCE): This error corresponds to an incorrect finite state machine with an extra or missing state (category 14). . Next state error (NSE): This error corresponds to incorrect next state function in a finite state machine (FSM) (category 14). Although, this extended set of error models increases the number of actual errors that can be modeled directly, we have found them to be too complex for practical use in manual or automated test generation. We observed that the more difficult actual errors are often composed of multiple basic errors, and that the component basic errors interact in such a way that a test to detect the actual error must be much more specific than a test to detect any of the component basic errors. Modeling these difficult composite errors directly is impractical as the number of error instances to be considered is too large and such composite modeled errors are too complex for automated test generation. However, as noted earlier, a good error model does not necessarily need to mimic actual errors accu- rately. What is required is that the error model necessitates the generation of these more specific tests. To be practical, the complexity of the new error models should be comparable to that of the basic error models. Furthermore the (unavoidable) increase in the number of error instances should be controlled to allow trade-offs between test generation effort and verification confidence. We found that these requirements can be combined by augmenting the basic error models with a condition. A conditional error (C,E) consists of a condition C and a basic error E; its interpretation is that E is only active when C is satisfied. In general, C is a predicate over the signals in the circuit during some time period. To limit the number of error instances, we restrict C to a conjunction of terms , where y i is a signal in the circuit and w i is a constant of the same bit-width as y i and whose value is either all-0s or all-1s. The number of terms (condition variables) appearing in C is said to be the order of (C,E). Specifically, we consider the following conditional error types: . Conditional single-stuck line (CSSLn) error of order n; . Conditional bus order error (CBOEn) of order n; . Conditional bus source error (CBSEn) of order n. When reduces to the basic error E from which it is derived. Higher-order conditional errors enable the generation of more specific tests, but lead to a greater test generation cost due to the larger number of error instances. For exam- ple, the number of CSSLn errors on a circuit with N signals is . Although the total set of all N signals we consider for each term in the condition can possibly be reduced, CSSLn errors where n > 2 are probably not practical. For gate-level circuits (where all signals are 1-bit), it can be shown that CSSL1 errors High-Level Design Verification of Microprocessors via Error Modeling - 11 cover the following basic error models: MSEs (excluding XOR and XNOR gates), missing 2-input gate errors, BSEs, single BCEs (excluding XOR and XNOR gates), and bus driver errors. That CSSL1 errors cover missing two-input gate errors can be seen as follows. Consider a two-input AND gate Y=AND(X1,X2) in the correct design; in the erroneous design, this gate is missing and net Y is identical to net X1. To expose this error we have to set X1 to 1, X2 to 0, and sensitize Y. Any test that detects the CSSL1 error, (X2=0, Y s-a-0) in the erroneous design, will also detect the missing gate error. The proof for other gate types is similar. Higher-order CSSLn errors improve coverage even further. 4. COVERAGE EVALUATION To show the effectiveness of a verification methodology, one could apply it and a competing methodology to an unverified design. The methodology that uncovers more (and hard- er) design errors in a fixed amount of time is more effective. However, for such a comparison to be practical, fast and efficient high-level test generation tools for our error models appear to be necessary. Although we have demonstrated such test generation in Section 3.2, it has yet to be automated. We therefore designed a controlled experiment that approximates the conditions of the original experiment, while avoiding the need for automated test generation. The experiment evaluates the effectiveness of our verification methodology when applied to two student-designed microprocessors. A block diagram of the experimental set-up is show in Figure 6. As design error models are used to guide test generation, the effectiveness is closely related to the synthetic error models used. To evaluate our methodology, a circuit is chosen for which design errors are to be systematically recorded during its design. Let D 0 be the final, presumably correct design. From the CVS revision database, the actual errors are extracted and converted such that they can be injected in the final design D 0 . In the evaluation phase, the design is restored to an (artificial) erroneous state D 1 by injecting a single actual error into the final design D 0 . This set-up approximates a realistic on-the-fly design verification scenario. The experiment answers the question: given D 1 , can the proposed methodology produce a test that determines D 1 to be erroneous? This is achieved by examining the actual error in D 1 , and determining if a modeled design error exists that is dominated by the actual error. Let D 2 be the design constructed by injecting the dominated modeled error in D 1 , and let M be the error model which defines the dominated modeled error. Such a dominated modeled error has the property that any test that detects the modeled error in D 2 will also detect the actual error in D 1 . Consequently, if we were to generate a complete test set for every error defined on D 1 by error model M, D 1 would be found erroneous by that test set. Error detection is determined as discussed earlier (see Section 1, Figure 2). Note that the concept of dominance in the context of design verification is slightly different than in physical fault testing. Unlike in the testing problem, we cannot remove the actual design error from D 1 before injecting the dominated modeled error. This distinction is important because generating a test for an error of omission, which is generally very hard, becomes easy if given instead of D 1 . The erroneous design D 1 considered in this experiment is somewhat artificial. In reality the design evolves over time as bugs are introduced and eliminated. Only at the very end of the design process, is the target circuit in a state where it differs from the final design D 0 in just a single design error. Prior to that time, the design may contain more than one Al-Asaad, J. P. Hayes, T. Mudge, and R. B. Brown design error. To the extent that the design errors are independent, it does not matter if we consider a single or multiple design errors at a time. Furthermore, our results are independent of the order in which one applies the generated tests. We implemented the preceding coverage-evaluation experiment for two small but representative designs: a simple microprocessor and a pipelined microprocessor. We present our results in the remainder of this section. 4.1 A simple microprocessor The Little Computer 2 (LC-2) [29] is a small microprocessor of conventional design used for teaching purposes at the University of Michigan. It has a representative set of 16 instructions which is a subset of the instruction sets of most current microprocessors. To serve as a test case for design verification, one of us designed behavioral and RTL synthesizable Verilog descriptions for the LC-2. The behavioral model (specification) of the LC-2 consists of 235 lines of behavioral Verilog code. The RTL design (implementation) consists of a datapath module described as an interconnection of library modules and a few custom modules, and a control module described as an FSM with five states. It comprises 921 lines of Verilog code, excluding the models for library modules such as adders, register files, etc. A gate-level model of the LC-2 can thus be obtained using logic synthesis tools. The design errors made during the design of the LC-2 were systematically recorded using our error collection system (Section 2). For each actual design error recorded, we derived the necessary conditions to detect it. An error is detected by an instruction sequence s if the external output signals of the behavioral and RTL models are distinguished by s. We found that some errors are undetectable since they do not affect the functionality of the microprocessor. The detection conditions are used to determine if a modeled error that is dominated by the actual error 6. Experiment to evaluate the proposed design verification methodology. Simulate Simulate Actual error database Debug by Design error collection Test for modeled error Evaluation of verification methodology Expose modeled error Expose actual error Design and debugging process Design Inject single actual error Inject modeled error Design error model designer Actual error Modeled error revisions High-Level Design Verification of Microprocessors via Error Modeling - 13 can be found. An example where we were able to do that is shown in Figure 7. The error is a BSE on data input D 1 of the multiplexer attached to the program counter PC. Testing for detect the BSE since the outputs of PC and its incrementer are always different, i.e., the error is always activated, so testing for this SSL error will propagate the signal on D 1 to a primary output of the microprocessor. A case where we were not able to find a modeled error dominated by the actual error is shown in Figure 8. The error occurs where a signal is assigned a value independent of any condition. However, the correct implementation requires an if-then-else construct to control the signal assignment. To activate this error, we need to set ir_out[15:12] == 4'b1101, ir_out[8:6] - 3'b111, and refers to the contents of the register i in the register file. An instruction sequence that detects this error is shown in Figure 8. We analyzed the actual design errors in both the behavioral and RTL designs of the LC- 2, and the results are summarized in Table 2. A total of 20 design errors were made during the design, of which four errors are easily detected by the Verilog simulator and/or logic synthesis tools and two are undetectable. The actual design errors are grouped by cate- gory; the numbers in parentheses refer to the corresponding category in Table 1. The columns in the table give the type of the simplest dominated modeled error corresponding to each actual error. For example, among the 4 remaining wrong-signal-source errors, 2 dom- 7. An example of an actual design error that is dominated by an SSL error. design Correct design Incrementer Incrementer Mux // Instruction decoding // Decoding of register file inputs // 1- Decoding of CORRECT CODE: if (ir_out[15:12] == 4'b1101) else ERRONEOUS CODE: 8. An example of an actual design error for which no dominated modeled error was found, and an instruction sequence that detects the actual error. // Instruction sequence @3000 main: JSR sub0 sub0: // After execution of instructions correct design Design error Test sequence 14 - D. Van Campenhout, H. Al-Asaad, J. P. Hayes, T. Mudge, and R. B. Brown inate an SSL error and 2 dominate a BSE error. We can infer from Table 2 that most errors are detected by tests for SSL errors or BSEs. About 75% of the actual errors in the LC-2 design can be detected after simulation with tests for SSL errors and BSEs. The coverage increases to 90% if tests for CSSL1 are added. 4.2 A pipelined microprocessor Our second design case study considers the well-known DLX microprocessor [19], which has more of the features found in contemporary microprocessors. The particular DLX version considered is a student-written design that implements 44 instructions, has a five-stage pipeline and branch prediction logic, and consists of 1552 lines of structural Verilog code, excluding the models for library modules such as adders, registerfiles, etc. The design errors committed by the student during the design process were systematically recorded using our error collection system. For each actual design error we painstakingly derived the requirements to detect it. detection was determined with respect to one of two reference models (specifica- tions). The first reference model is an ISA model, and as such is not cycle-accurate: only the changes made to the ISA-visible part of the machine state, that is, to the register file and memory, can be compared. The second reference model contains information about the microarchitecture of the implementation and gives a cycle-accurate view of the ISA- visible part of the machine state (including the program counter). We determined for each actual error whether it is detectable or not with respect to each reference model. Errors undetectable with respect to both reference models may arise from the following two rea- sons: (1) Designers sometimes make changes to don't care features, and log them as errors. This happens because designers can have a more detailed specification (design in mind than that actually specified. (2) Inaccuracies can occur when fixing an error requires multiple revisions. We analyzed the detection requirements of each actual error and constructed a modeled 2. Actual design errors and the corresponding dominated modeled errors for LC-2. Actual errors Corresponding dominated modeled errors Category Total Easily detected Undetec- table SSL BSE CSSL1 Un- known Wrong signal source (1) 4 Expression error (7) 4 Bit width error Missing assignment Wrong constant Unused signal Wrong module (5) 1 Always statement Total High-Level Design Verification of Microprocessors via Error Modeling - 15 error dominated by the actual error, wherever possible. One actual error involved multiple signal source errors, and is shown in Figure 9. Also shown are the truth tables for the immediately affected signals; differing entries are shaded. Error detection via fanout Y1 requires setting sensitizing Y1. However, the combination not achievable and thus error detection via Y1 is not possible. Detection via fanout Y2 or Y3 requires setting sensitizing Y2 or Y3. However, blocks error propagation via Y2 further downstream. Hence, the error detection requirements are: sensitizing Y3. Now consider the modeled error . Activation of E 1 in D1 requires sensitizing Y1, Y2 or Y3. As mentioned before, blocks error propagation via Y2. But as E 1 can be exposed via Y1 without sensitizing Y3, E 1 is not dominated by the given actual error. To ensure detection of the actual error, we can condition S0 s-a-0 such that sensitization of Y3 is required. The design contains a signal jump_to_reg_instr that, when set to 1, blocks sensitization of Y1, but allows sensitization of Y3. Hence the CSSL1 error dominated by the actual error. The results of this experiment are summarized in Table 3. A total of 39 design errors were recorded by the designer. The actual design errors are grouped by category; the numbers in parentheses refer again to Table 1. The correspondence between the categories is imprecise, because of inconsistencies in the way in which different student designers classified their errors. Also, some errors in Table 3 are assigned to a more specific category than in Table 1, to highlight their correlation with the errors they dominate. 'Missing mod- ule' and `wrong signal source' errors account for more than half of all errors. The column headed 'ISA' indicates how many errors are detectable with respect to the ISA-model; 'ISAb' lists the number of errors only detectable with respect to the micro-architectural reference model. The sum of 'ISA' and `ISAb' does not always add up the number given 9. Example of an actual design error in our DLX implementation. D. Van Campenhout, H. Al-Asaad, J. P. Hayes, T. Mudge, and R. B. Brown in 'Total'; the difference corresponds to actual errors that are not undetectable with respect to either reference model. The remaining columns give the type of the simplest dominated modeled error corresponding to each actual error. Among the 10 detectable 'missing mod- ule(s)' errors, 2 dominate an SSL error, 6 dominate a CSSL1 error, and one dominates a CBOE; for the remaining one, we were not able to find a dominated modeled error. A conservative measure of the overall effectiveness of our verification approach is given by the coverage of actual design errors by complete test sets for modeled errors. From Table 3 it can be concluded that for this experiment, any complete test set for the inverter insertion errors (INV) also detects at least 21% of the (detectable) actual design errors; any complete test set for the INV and SSL errors covers at least 52% of the actual design errors; if a complete test set for all INV, SSL, BSE, CSSL1 and CBOE is used, at least 94% of the actual design errors will be detected. 5. DISCUSSION The preceding experiments indicate that a high coverage of actual design errors can be obtained by complete test sets for a limited number of modeled error types, such as those defined by our basic and conditional error models. Thus our methodology can be used to construct focused test sets aimed at detecting a broad range of actual design bugs. More impor- tantly, perhaps, it also supports an incremental design verification process that can be implemented as follows: First, generate tests for SSL errors. Then generate tests for other basic error types such as BSEs. Finally, generate tests for conditional errors. As the number of SSL errors in a circuit is linear in the number of signals, complete test sets for SSL errors can be relatively small. In our experiments such test sets already detect at least half of the actual errors. To improve coverage of actual design errors and hence increase the confidence in the design, an error model with a quadratic number of error instances, such as BSE and CSSL1, can be used to guide test generation. The conditional error models proved to be especially useful for detecting actual errors that involve missing logic. Most 'missing module(s)' and `missing input(s)' in Table 3 3. Actual design errors and the corresponding dominated modeled errors for DLX. Actual errors Corresponding dominated modeled errors Category ISA ISAb Total INV SSL BSE CSSL1 CBOE CSSL2 Un- known Missing module (2) 8 2 14 Wrong singal source (1) 9 2 Complex Inversion Missing input Unconnected input Missing minterm (2) Extra input (2) Total High-Level Design Verification of Microprocessors via Error Modeling - 17 cannot be covered when only the basic errors are targeted. However, all but one of them is covered when CSSL1 and CBOE errors are targeted as well. The same observation applies to the 'missing assignment(s)' errors in Table 2. The designs used in the experiments are small, but appear representative of real industrial designs. An important benefit of such small-scale designs is that they allow us to analyze each actual design error in detail. The coverage results obtained strongly demonstrate the effectiveness of our model-based verification methodology. Furthermore the analysis and conclusions are independent of the manner of test generation. Nevertheless, further validation of the methodology using industrial-size designs is desirable, and will become more practical when CAD support for design error test generation becomes available. models of the kind introduced here can also be used to compute metrics to assess the quality of a given verification test set. For example, full coverage of basic (uncondi- tional) errors provides one level of confidence in the design, coverage of conditional errors of order provides another, higher confidence level. Such metrics can also be used to compare test sets and to direct further test generation. We envision the proposed methodology eventually being deployed as suggested in Figure 2. Given an unverified design and its specification, tests targeted at modeled design errors are automatically generated and applied to the specification and the implementation. When a discrepancy is encountered, the designer is informed and perhaps given guidance on diagnosing and fixing the error. ACKNOWLEDGMENTS We thank Steve Raasch and Jonathan Hauke for their help in the design error collection process. We further thank Matt Postiff for his helpful comments. The research discussed in this paper is supported by DARPA under Contract No. DABT63-96-C-0074. The results presented herein do not necessarily reflect the position or the policy of the U.S. Government. --R "Logic design verification via test generation," "Verification of the IBM RISC System/6000 by dynamic biased pseudo-random test program generator" "Design verification via simulation and automatic test pattern gen- eration" "High-level design verification of microprocessors via error modeling," "From specification validation to hardware testing: A unified method" "High-level test generation using bus faults," New York "A neutral netlist of 10 combinational benchmark circuits and a target translator in fortran" "Complementary GaAs technology for a GHz microprocessor" Cadence Design Systems Inc. "Functional verification methodology of Chameleon processor" Version Management with CVS "AVPGEN - a test generator for architecture verification" "Latent design faults in the development of the Multiflow TRACE/200" "Hints on test data selection: Help for the practicing programmer" "Observability-based code coverage metric for functional simulation" "Hardware emulation for functional verification of K5" "High-level test generation using physically-induced faults" Computer Architecture: A Quantitative Approach "Code generation and analysis for the functional verification of microprocessors" "Pentium Processor Specification Update," "I'm done simulating; Now what? Verification coverage analysis and correctness checking of the DECchip 21164 Alpha microprocessor" high-level Verilog description" "A Fortran language system for mutation-based software testing" "Prototyping the M68060 for concurrent verification" MIPS Technologies Inc. "An experimental determination of sufficient mutant operators" "Finite state machine trace analysis program" The TTL Logic Data Book "Mutation testing - its origin and evolution" Formal Verification of Hardware Design --TR Computer architecture: a quantitative approach Software testing techniques (2nd ed.) A Fortran language system for mutation-based software testing Verification of the IBM RISC System/6000 by a dynamic biased pseudo-random test program generator AVPGENMYAMPERSANDmdash;a test generator for architecture verification Design verification via simulation and automatic test pattern generation An experimental determination of sufficient mutant operators Code generation and analysis for the functional verification of micro processors Hardware emulation for functional verification of K5 I''m done simulating; now what? Verification coverage analysis and correctness checking of the DEC chip 21164 Alpha microprocessor Functional verification methodology of Chameleon processor An observability-based code coverage metric for functional simulation Formal Verification of Hardware Design Prototyping the M68060 for Concurrent Verification From Specification Validation to Hardware Testing High-level test generation using physically-induced faults --CTR David Van Campenhout , Trevor Mudge , John P. Hayes, Collection and Analysis of Microprocessor Design Errors, IEEE Design & Test, v.17 n.4, p.51-60, October 2000 Katarzyna Radecka , Zeljko Zilic, Identifying Redundant Wire Replacements for Synthesis and Verification, Proceedings of the 2002 conference on Asia South Pacific design automation/VLSI Design, p.517, January 07-11, 2002 Tao Lv , Jian-Ping Fan , Xiao-Wei Li , Ling-Yi Liu, Observability Statement Coverage Based on Dynamic Factored Use-Definition Chains for Functional Verification, Journal of Electronic Testing: Theory and Applications, v.22 n.3, p.273-285, June 2006 Katarzyna Radecka , Zeljko Zilic, Design Verification by Test Vectors and Arithmetic Transform Universal Test Set, IEEE Transactions on Computers, v.53 n.5, p.628-640, May 2004 Wei Lu , Xiu-Tao Yang , Tao Lv , Xiao-Wei Li, An efficient evaluation and vector generation method for observability-enhanced statement coverage, Journal of Computer Science and Technology, v.20 n.6, p.875-884, November 2005 Anand L. 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design errors;error modeling;design verification

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Extracting Hidden Context.

Concept drift due to hidden changes in context complicates learning in many domains including financial prediction, medical diagnosis, and communication network performance. Existing machine learning approaches to this problem use an incremental learning, on-line paradigm. Batch, off-line learners tend to be ineffective in domains with hidden changes in context as they assume that the training set is homogeneous. An off-line, meta-learning approach for the identification of hidden context is presented. The new approach uses an existing batch learner and the process of {\it contextual clustering} to identify stable hidden contexts and the associated context specific, locally stable concepts. The approach is broadly applicable to the extraction of context reflected in time and spatial attributes. Several algorithms for the approach are presented and evaluated. A successful application of the approach to a complex flight simulator control task is also presented.

Introduction In real world machine learning problems, there can be important properties of the domain that are hidden from view. Furthermore, these hidden properties may change over time. Machine learning tools applied to such domains must not only be able to produce classifiers from the available data, but must be able to detect the effect of changes in the hidden properties. For example, in finance, a successful stock buying strategy can change dramatically in response to interest rate changes, world events, or with the season. As a result, concepts learnt at one time can subsequently become inaccurate. Concept drift occurs with changes in the context surrounding observations. Hidden changes in context cause problems for any machine learning approach that assumes concept stability. In many domains, hidden contexts can be expected to recur. These domains include: financial prediction, dynamic control and other commercial data mining applications. Recurring contexts may be due to cyclic phenomena, such as seasons of the year or may be associated with irregular phenomena, such as inflation rates or market mood. Machine learning systems applied in domains with hidden changes in context have tended to be incremental or on-line systems, where the concept definition is updated as new labeled observations are processed (Schlimmer & Granger, 1986). Adaptation to new domains is generally achieved by decaying the importance of older instances. Widmer and Kubat's on-line system, Flora3 (Widmer & Kubat, 1993), exploits recurring hidden context. As the system traverses the sequence of input data, it stores concepts that appear to be stable over some interval of time. These stable concepts can be retrieved allowing the algorithm to adapt quickly when the observed domain changes to a context previously encountered. Stable concepts can also be identified off-line using batch learning algorithms. Financial institutions, manufacturing facilities, government departments, etc., all store large amounts of historical data. Such data can be analyzed, off-line, to discover regularities. Patterns in the data, however, may be affected by context changes without any records of the context being maintained. To handle these sit- uations, the batch learner must be augmented to detect hidden changes in context. Missing context is often reflected in the temporal proximity of events. For exam- ple, there may be days in which all customers buy chocolates. The hidden context in this case might be that a public holiday is due in the following week. Hidden context can also be distributed over a non-temporal dimension, making it completely transparent to on-line learners. For example, in remote sensing, the task of learning to classify trees by species may be affected by the surrounding forest type. If the forest type is not available to the learning system, it forms a hidden context, distributed by geographic region rather than time. Off-line methods for finding stable concepts can be applied to these domains. For simplicity, this article retains the convention of organizing hidden context over time but the methods presented generalize to properties other than time. Each stable concept is associated with one or more intervals in time. The shift from one stable concept to another represents a change in context. Thus each interval can be identified with a particular context. This presents the opportunity to build models of the hidden context. Such a model may be desirable for explanatory purposes to understand a domain or it may be incorporated into an on-line predictive model. A model may also be used to identify a new attribute that correlates with the hidden context. In this paper, we present Splice, an off-line meta-learning system for context-sensitive learning. Splice is designed to identify stable concepts during supervised learning in domains with hidden changes in context. We begin by reviewing related work on machine learning in context-sensitive do- mains. This is followed by a description of the Splice methodology. An initial implementation of Splice, Splice-1, was previously shown to improve on a standard induction method in simple domains. We briefly discuss this work before presenting an improved algorithm, Splice-2. Splice-2 is shown to be superior to in more complex domains. The Splice-2 evaluation concludes with an application to a complex control task. 2. Background On-line learning methods for domains with hidden changes in context adapt to new contexts by decaying the importance of older instances. Stagger (Schlimmer and Granger, 1996), was the first reported machine learning system that dealt with hidden changes in context. This system dealt with changes in context by discarding any concepts that fell below a threshold accuracy. Splice is most related to the Flora (Widmer & Kubat, 1996) family of on-line learners. These adapt to hidden changes in context by updating the current concept to match a window of recent instances. Rapid adaptation to changes in context is assured by altering the window size in response to shifts in prediction accuracy and concept complexity. One version, Flora3, (Widmer & Kubat, 1993) adapts to domains with recurring hidden context by storing stable concepts, these can be re-used whenever context change is suspected. When a concept is re-used, it is first updated to match examples in the current window. This allows Flora3 to deal with discrepancies between the recalled concept and the actual situation. Rather than an adjunct to on-line learning, Splice makes the strategy of storing stable concepts the primary focus of an off-line learning approach. Machine learning with an explicit window of recent instances, as used in Flora, was first presented by Kubat (1989) and has been used in many other on-line systems dealing with hidden changes in context. The approach has been used for supervised learning (Kubat & Widmer, 1995), and unsupervised learning (Kilander Jansson, 1993). It is also used to adapt batch learners for on-line learning tasks by repeatedly learning from a window of recent instances (Harries & Horn, 1995; 1989). The use of a window can also be made sensitive to changes in the distribution of instances. Salganicoff (1993) replaces the first in, first out updating method by discarding older examples only when a new item appears in a similar region of attribute space. Most batch machine learning methods assume that the training items are independent and unordered. They also assume that all available information is directly represented in the attributes provided. As a result, batch learners generally treat hidden changes in context as noise. For example, Sammut, Hurst, Kedzier and Michie (1992), report on learning to pilot an aircraft in a flight simulator. They note that a successful flight could not be achieved without explicitly dividing the flight into stages. In this case, there were known changes of context and so, learning could be broken into several sub-tasks. Within each stage of the flight, the control strategy, that is, the concept to be learnt, was stable. Splice applies the assumption that concepts are likely to be stable for some period of time to the problem of detecting stable concepts and extracting hidden context. 3. SPLICE Splice's input is a sequence of training examples, each consisting of a feature vector and a known classification. The data are ordered over time and may contain hidden 4changes of context. From this data, Splice attempts to learn a set of stable concepts, each associated with a different hidden context. Since contexts can recur, several disjoint intervals of the data set may be associated with the same concept. On-line learners for domains with hidden context assume that a concept will be stable over some interval of time. Splice also uses this assumption for batch learning. Hence, sequences of examples in the data set are combined into intervals if they appear to belong to the same context. Splice then attempts to cluster similar intervals by applying the notion that similarity of context is reflected by the degree to which intervals are well classified by the same concept. This is called contextual clustering. Informally, a stable concept is an expression that holds true for some period of time. The difficulty in finding a stable concept is in determining how long "some period" should be. Clearly, many concepts may be true for very short periods. Splice uses a heuristic to divide the data stream into a minimal number of partitions (contextual clusters), which may contain disjoint intervals of the dataset, so that a stable concept created from one contextual cluster will poorly classify examples in all other contextual clusters. In a sense, the stable concept will be the most specific concept describing the contextual cluster. A more rigorous method for comparing different sets of contextual clusters might use a Minimum Description Length (MDL) measure (Rissanen, 1983). The MDL principle states that the best theory for a given concept should minimize the amount of information that needs be sent from a sender to a receiver so that the receiver can correctly classify items in a shared dataset. In this case, the information to be sent would include stable concepts, a context switching method and a list of exceptions. A good set of contextual clusters should result in stable concepts that give a shorter description length for describing the data than would a single concept. An optimal set of contextual clusters should achieve the minimum description length possible. A brute force approach to finding a set of clusters that satisfy the MDL measure would be to consider all possible combinations of contextual clusters in the dataset, then select the combination with the minimum description length. Clearly, this is impractical. Manganaris (1996) applies the minimum description length heuristic to the creation of a piecewise polynomial function from a series of numbers (method adapted from Pednault (1989). With dynamic programming the space of possible partitions can be searched in O(n 2 ) in time. Adapting this method for Splice would give time complexity of O(n 4 ). Therefore, Splice uses a heuristic approach to find stable concepts that are "good enough". The Splice algorithm is a meta-learning algorithm. Concepts are not induced directly, but by application of an existing batch leaner. In this study, we use Quinlan's C4.5 (Quinlan, 1993), but the Splice methodology could be implemented using other propositional learning systems. C4.5 is used without modification. Furthermore, since noise is dealt with by C4.5, Splice contains no explicit noise handling mechanism. Unusual levels of noise can be dealt with by altering the C4.5 parameters. The main purpose of this paper is to present the Splice-2 algorithm. However, we first briefly describe it's predecessor, Splice-1 (Harries & Horn, in press) and its shortcomings to motivate the development of the Splice-2 algorithm. 3.1. SPLICE-1 first uses a heuristic to identify likely context boundaries. Once the data has been partitioned on these boundaries, the partitions are combined according to their similarity of context. Stable concepts are then induced from the resulting contextual clusters. Details of the Splice-1 algorithm have previously been reported by Harries and Horn (in press) so we only give a brief overview in this section. To begin, each example is time-stamped to give its position in the sequence of training data. Thus, time forms a continuous attribute in which changes of context can be expressed. For example, the hidden context, interest rate, might change at time=99. C4.5 is then used to induce a decision tree from the whole training set. Each node of the tree contains a test on an attribute. Any test on the special attribute, time, is interpreted as indicating a possible change of context. For example, Table 1 shows a simple decision tree that might be used for stock market investment. This tree includes a test on time, which suggests that a change in context may have occurred at Time=1995. Splice-1 then uses 1995 as boundary to partition the data set. We assume that each interval, as defined by all such partitions, can be identified with a stable concept. Table 1. Sample decision tree in a domain with hidden changes in context. Attribute Attribute Attribute Attribute Time ?= 1995 If a stable concept induced from one interval accurately classifies the examples in another interval, we assume that both intervals have similar contexts. The degree of accuracy provides a continuous measure of the degree of similarity. Intervals are grouped by contextual similarity. When adjacent intervals are combined, a larger context is identified. When disjunctive sub-sets are combined, a recurring context is identified. C4.5 is applied again to the resulting contextual clusters to produce the final stable concepts. 3.2. SPLICE-1 prediction Harries and Horn (in press) have shown that Splice-1 can build more accurate classifiers than a standard induction algorithm in sample domains with hidden Splice Machine Learner Dynamic Concept Switching Training stable concepts Figure 1. Splice on-line prediction changes in context. We summarize these results and provide a comparison with the on-line method Flora3. In this task, Splice-1 is used to produce a set of stable concepts, which are then applied to an on-line prediction task. Figure 1 shows this process schematically. On-line classification is achieved by switching between stable concepts according to the current context. 3.2.1. STAGGER data set The data sets in the following experiments are based on those used to evaluate Stagger (Schlimmer & Granger, 1986) and subsequently used to evaluate Flora (Widmer & Kubat, 1996). While our approach is substantially different, use of the same data set allows some comparison of results. A program was used to generate data sets. This allows us to control the recurrence of contexts and other factors such as noise 1 and duration of contexts. The task has four attributes, time, size, color and shape. Time is treated as a continuous attribute. Size has three possible values: small, medium and large. Color has three possible values: red, green and blue. Shape also has three possible values: circular, triangular, and square. The program randomly generates a sequence of examples from the above attribute space. Each example is given a unique time stamp and a boolean classification based upon one of three target concepts. The target concepts are: 1. 2. 3. Artificial contexts were created by fixing the target concepts to one of the above Stagger concepts for preset intervals of the data series. 3.2.2. On-line prediction This experiment compares the accuracy of Splice-1 to C4.5 when trained on a data set containing changes in a hidden context. In order to demonstrate that the Splice approach is valid for on-line classification tasks, we show a sample prediction problem in which Splice-1 was used off-line to generate a set of stable concepts from training data. The same training data was used to generate a single concept using C4.5, again off-line. After training, the resulting concepts were applied to a simulated on-line prediction task. C4.5 provides a baseline performance for this task and was trained without the time attribute. C4.5 benefits from the omission of this attribute as the values for time in the training set do not repeat in the test set. Even so, this comparison is not altogether fair on C4.5, as it was not designed for use in domains with hidden changes in context. The training set consisted of concept (1) for 50 instances, (2) for 50 instances, and (3) for 50 instances. The test set consisted of concepts (1) for 50 instances, (2) for 50 instances, (3) for 50 instances, and repeated (1) for 50 instances, (2) for 50 instances, and (3) for 50 instances. To apply the stable concepts identified by Splice for prediction, it was necessary to devise a method for selecting relevant stable concepts. This is not a trivial problem. Hence, for the purposes of this experiment we chose a simple voting method. With each new example, the classification accuracy of each stable concept over the last five examples was calculated. The most accurate concept was then used to classify the new example. Any ties in accuracy were resolved by random selection. The first case was classified by a randomly selected stable concept. Figure 2 shows the number of correct classifications for each item in the test set for both Splice-1 and C4.5 over 100 randomly generated training and test sets. Figure 2 shows that Splice-1 successfully identified the stable concepts from the training set and that the correct concept can be successfully selected for prediction in better than 95% of cases. The extreme dips in accuracy when contexts change are an effect of the method used to select stable concepts. C4.5 performs relatively well on concept 2 with an accuracy of approximately 70% but on concepts 1 and 3, it correctly classifies between 50% and 60% of cases. As noise increases, the performance of Splice-1 gradually declines. At 30% noise (Harries & Horn, in press), the worst result achieved by Splice-1 is an 85% classification accuracy on concept 2. C4.5, on the other hand, still classifies with approximately the same accuracy as it achieved in Figure 2. This task is similar to the on-line learning task attempted using Flora (Widmer Kubat, 1996) and Stagger (Schlimmer & Granger, 1986). The combination of Splice-1 with a simple strategy for selection of the current stable concept is effective on a simple context sensitive prediction task. As the selection mechanism assumes that at least one of the stable concepts will be correct, Splice-1 almost immediately moves to its maximum accuracy on each new stable concept. For a similar problem, the Flora family (Widmer & Kubat, 1996) (in particular Flora3, the learner designed to exploit recurring context) appear to reach much the same level of accuracy as Splice-1, although as an on-line learning method, Flora requires some time to fully reflect changes in context. correct item number Splice Figure 2. On-line prediction task. Compares Splice-1 stable concepts with a single C4.5 concept. This comparison is problematic for a number of reasons. Splice-1 has the advantage of first seeing a training set containing 50 instances of each context before beginning to classify. Furthermore, the assumption that all possible contexts have been seen in the training set is correct for this task. On-line learners have the advantage of continuous feedback with an unconstrained updating of concepts. Splice-1 does have feedback, but is constrained to select only from those stable concepts learnt from the training data. When Splice-1 has not learnt a stable concept, there is no second chance. For more complex domains, it could be beneficial to use a combination of Splice-1 and an adaptive, on-line learner. 4. SPLICE-2 uses the assumption that splits on time resulting from a run of C4.5 will accurately reflect changes in context. This is not always true. Splice-2 was devised to reduce the reliance on initial partitioning. Like Splice-1, Splice-2 is a "meta-learner". In these experiments we again use C4.5 (Quinlan, 1993) as the induction tool. The Splice-2 algorithm is detailed in Figure 2. The two stages of the algorithm are discussed in turn. 4.1. Stage 1: Partition dataset begins by guessing an initial partitioning of the data and subsequent stages refine the initial guess. Any of three methods may be used for the initial guess: ffl Random partitioning. Randomly divide the data set into a fixed number of partitions. ffl Partitioning by C4.5. As used in Splice-1. Each test on time, found when C4.5 is run on the entire data set, is used as an initial partition. Table 2. The Splice-2 Algorithm Ordered data set, Window size parameter.) Partition Dataset - Partition the dataset over time using either: A preset number of random splits. C4.5 as per Splice-1. Prior domain knowledge. - The identified partitions form the initial contextual clusters. - C4.5 is applied to the initial contextual clusters to produce the initial interim concepts. ffl Stage 2: Contextual Clustering Each combination of interim concept and item in the original data set is allocated a score based upon the total accuracy of that concept on items in a fixed size window over time surrounding the item. - Cluster the original data set items that share maximum scores with the same concept. These clusters form the new set of contextual clusters. - C4.5 is used to create a new set of interim concepts from the new contextual clusters. Stage 2 is repeated until the interim concepts do not change or until a fixed number of iterations are completed. - The last iteration provides the set of stable concepts. dataset order initial contextual clusters initial concepts Figure 3. Splice-2: stage 1 ffl Prior Domain Knowledge. In some domains, prior knowledge is available about likely stable concepts. We denote the version of Splice-2 using random partitioning as Splice-2R, the version using C4.5 partitioning as Splice-2C, and the version using prior domain knowledge as Splice-2P. Once the dataset has been partitioned, each interval of the dataset is stored as an initial contextual cluster. C4.5 is applied to each cluster to produce a decision tree known as an interim concept, see Figure 3. 4.2. Stage 2: Contextual clustering Stage 2 iteratively refines the contextual clusters generated in stage 1. With each iteration, a new set of contextual clusters is created in an attempt to better identify stable concepts in the data set. This stage proceeds by testing each interim concept for classification accuracy against all items in the original data set. A score is computed for each pair of concept and item number. This score is based upon the number of correct classifications achieved in a window surrounding the item (see Figure 4). The window is designed to capture the notion that a context is likely to be stable for some period of time. On-line learning systems apply this notion implicitly by using a window of recent instances. Many of these systems use a fixed window size, generally chosen to be the size of the minimum context duration expected. The window in Splice has the same function as in an on-line method, namely, to identify a single context. Ideally, the window size should be dynamic, as in Flora, allowing the window to adjust Figure 4. Splice-2: stage 2. Using interim concept accuracy over a window to capture context. for different concepts. For computational efficiency and simplicity, we use a fixed sized window. The context of an item can be represented by a concept that correctly classifies all items within the window surrounding that item. At present the window size is set to 20 by default. This window size was chosen so as to bias contexts identified to be of more than 20 instances duration. This was considered to be the shortest context that would be valuable during on-line classification. The window size can be altered for different domains. We define W ij to be the score for concept j when applied to a window centered on example i. Correct jm (1) where: correctly classifies example m the window size The current contextual clusters, from stage one or a previous iteration of stage two, are discarded. New contextual clusters are then created for each interim concept j. Once all scores are computed, each item, i, is allocated to a contextual cluster associated with the interim concept, j, that maximizes W ij over all interim concepts. These interim concepts are then discarded. C4.5 is applied to each new contextual cluster to learn a new set of interim concepts (Figure 5). The contextual clustering process iterates until either a fixed number of repetitions is completed or until the interim concepts do not change from one iteration to the next. The last iteration of this stage provides a set of stable concepts. The final contextual clusters give the intervals of time for which different contexts are active. 4.2.1. Alternate weight In domains with a strong bias toward a single class, such as 90% class A and 10% class B, well represented classes can dominate the contextual clustering process. This can lead to clusters that are formed by combining l clusters' concepts' Figure 5. Splice-2: stage 2. Create interim concepts. contexts with similar classifications for well represented classes and quite dissimilar classification over poorly represented classes. This is problematic in domains where the correct classification of rare classes is important, as in "learning to fly" (see Section 6). For such domains, the can be altered to give an equal importance to accuracy on all classes in a window while ignoring the relative representations of the different classes. For a given item, i, and an interim concept, j, the new sums over all possible classifications the proportion of items with a given class correctly classified. The (2) where: C is the number of classes c m is the class number of example m correctly classifies example m the window size 4.3. SPLICE-2 walk through In this section we present a walk through of the Splice-2R algorithm, using a simple dataset with recurring context. The dataset consists of randomly generated Stagger instances classified according to the following pattern of context: 3 repetitions of the following structure (concept (1) for instances, concept (2) for instances, and concept (3) for was applied to the dataset. begins by randomly partitioning the dataset into four periods. Each partition is then labeled as a contextual cluster. Figure 6 shows the instances r item no Figure 6. Initial contextual clusters. Created by randomly partitioning the dataset. associated with each contextual cluster as drawn from the original dataset. C4.5 is applied to each of these clusters, CC i , to induce interim concepts, IC i . Table 3 shows the induced concepts. Of these, only IC 4 relates closely to a target concept. Table 3. Initial interim concepts. no (34/7.1) Color in -red,green-: yes (70.0/27.3) IC3 Color in -red,blue-: - Size in -medium,large-: yes (14.0/7.8) IC4 Size in -medium,large-: - Shape in -square,triangular-: - Color in -red,blue-: no (33.0/16.5) calculated for all combinations of item, i, and interim concept, j. Table 4 shows the calculated W ij scores for a fragment of the dataset. Each figure in this table represents the number of items in a window surrounding the item number classified correctly by a given interim concept, IC j . For instance, the interim concept, IC 3 , classifies 12 items correctly in a window surrounding item number Each item, i, is then allocated to a new contextual cluster based upon the interim concept, IC j , that yielded the highest score, W ij . This is illustrated in Table 4 where the highest score in each column is italicized. At the base of the table we note the new contextual cluster to which each item is to be allocated. Each new contextual cluster is associated with a single interim concept. For example, interim concept IC 1 gives the highest value for item 80, so that item is allocated to a new contextual cluster item no Figure 7. Contextual clusters0. Created by the first iteration of the contextual clustering process.24 contextual cluster item no Figure 8. Contextual clusters000. Created by the third and final iteration of contextual clustering. contextual cluster, CC 1 0. This table contains items allocated to more than one contextual cluster, and implies a change of context around item 182. Table scores for the initial interim concepts on all items in the training set. Concepts Item Number (i) Allocate to The new contextual clusters, CC i 0, are expected to improve on the previous clusters by better approximating the hidden contexts. Figure 7 shows the distribution of the new contextual clusters, CC i 0. From these clusters, a new set of interim concepts, IC i 0 are induced. Contextual clustering iterates until either a fixed number of repetitions is completed or until the interim concepts do not change from one iteration to the next. Figure 8 shows the contextual clusters after two more iterations of the clustering stage. Each remaining contextual cluster corresponds with a hidden context in the original training set. At this point, the final interim concepts, IC i 000, are renamed as stable concepts, SC i . Stable concept one, SC 1 000 is the first target concept. Stable concept three, is the third target concept. Stable concept four, SC 4 000, is the second target concept. Contextual cluster two, CC 2 000, did not contain any items so was not used to induce a stable concept. 5. Experimental comparison of SPLICE-1 and SPLICE-2 This section describes two experiments comparing the performance of Splice-2 with Splice-1. The first is a comparison of clustering performance across a range of duration and noise levels with the number of hidden changes in context fixed. The second compares the performance of the systems across a range of noise and hidden context changes with duration fixed. In these experiments performance is determined by checking if the concepts induced by a system agree with the original concepts used in generating the data. Both these experiments use Stagger data, as described in the prior experiment. However, unlike the prior experiment only training data is required as we are assessing performance against the original concepts. 5.1. Contextual clustering: SPLICE-1 'vs' SPLICE-2 This experiment was designed to compare the clustering performance of the two systems. Splice-1 has been shown to correctly induce the three Stagger concepts under a range of noise and context duration conditions (Harries & Horn, in press). This experiment complicates the task of identifying concepts by including more context changes. In order to compare the clustering stage alone, Splice-2C was used to represent Splice-2 performance. This ensured that the initial partitioning used in both versions of Splice was identical. Both versions of splice were trained on independently generated data sets. Each data set consisted of examples classified according to the following pattern of con- repetitions of the following structure (concept (1) for D instances; concept (2) for D instances; and concept (3) for D instances), where the duration D ranged from 10 to 100 and noise ranged from 0% to 30%. This gave a total of 14 context changes in each training set. Splice-2 was run with the default window size of 20. The number of iterations of contextual clustering were set at three. C4.5 was run with default parameters and with sub-setting on. The stable concepts learnt by each system were then assessed for correctness against the original concepts. The results were averaged over 100 repetitions at each combination of noise and duration. They show the proportion of correct stable concept identifications found and the average number of incorrect stable concepts identified. Figure 9 shows the number of concepts correctly identified by Splice-1 and for a range of context durations and noise levels. The accuracy of both versions converges to the maximum number of concepts at the higher context durations for all levels of noise. Both versions show a graceful degradation of accuracy when noise is increased. Splice-2 recognizes more concepts at almost all levels of noise and concept duration. concepts correctly identified context duration no noise 10% noise 20% noise 30% noise0.51.52.510 20 concepts correctly identified context duration no noise 10% noise 20% noise 30% noise Figure 9. Concepts correctly identified by Splice-1 and Splice-2 when duration is changed2610 incorrect concepts context duration 0% noise 10% noise 20% noise 30% noise2610 incorrect concepts context duration 0% noise 10% noise 20% noise 30% noise Figure 10. Incorrect concepts identified by Splice-1 and Splice-2 Figure compares the number of incorrect concepts identified by Splice-1 and Splice-2. Both versions show a relatively stable level of performance at each level of noise. For all levels of noise, Splice-2 induces substantially fewer incorrect concepts than Splice-1. As both versions of Splice used the same initial partitioning, we conclude that the iterative refinement of contextual clusters, used by Splice-2, is responsible for the improvement on Splice-1. These results suggests that the Splice-2 clustering mechanism is better able to overcome the effects of frequent context changes and high levels of noise. In the next experiment we further investigate this hypothesis by fixing context duration and testing different levels of context change. 5.2. The effect of context repetition The previous experiment demonstrated that Splice-2 performs better than Splice- 1 with a fixed number of context changes. However the experiment provided little insight into the effect of different levels of context repetition. This experiment investigates this effect by varying the number of hidden context changes in the data. Once again we use Splice-2C to ensure the same partitioning used for Splice-1. We also examine the results achieved by Splice-2R. In this experiment each system was trained on an independent set of data that consisted of the following pattern of contexts: R repetitions of the structure (con- cept (1) for 50 instances; concept (2) for 50 instances; concept (3) for 50 instances) where R varies from one to five. The effects of noise were also evaluated with a range of noise from 0% to 40%. Splice-2 was run with the default window size of 20 with three iterations of contextual clustering. used partitions. The parameters for C4.5 and the performance measure used were the same as used in the prior experiment. Figure 11 shows the number of concepts correctly induced by both Splice-1 and Splice-2C for each combination of context repetition and noise. The results indicate that both systems achieve similar levels of performance with one context repetition. Splice-2C performs far better than Splice-1 for almost all levels of repetition greater than one. Comparing the shapes of the performance graphs for each system is interesting. shows an increasing level of performance across almost all levels of noise with an increase in repetition (or context change). On the other hand, Splice- 1 show an initial rise and subsequent decline in performance as the number of repetitions increases. The exception is at 0% noise, where both versions identify all three concepts with repetition levels of three and more. Figure 12 shows the number of correct Stagger concepts identified by Splice- 2R. The results shows a rise in recognition accuracy as repetitions increase (up to the maximum of 3 concepts recognized) for all noise levels. The number of concepts recognized is similar to those in Figure 11 for Splice-2C. concepts correctly identified context repetition 0% noise 10% noise 20% noise 30% noise concepts correctly identified context repetition 0% noise 10% noise 20% noise 30% noise 40% noise Figure 11. Concepts correctly identified by Splice-1 and Splice-2C across a range of context concepts correctly identified context repetition 0% noise 10% noise 20% noise 30% noise 40% noise Figure 12. Splice-2R concept identification The similarity of results for Splice-2C and Splice-2R shows that, for this do- main, C4.5 partitioning provides no benefit over the use of random partitioning for The results of this experiment indicate that Splice-2C is an improvement on as it improves concept recognition in response to increasing levels of context repetition. The performance of Splice-1 degrades with increased levels of context changes. The inability of Splice-1 to cope with high levels of context change is probably due to a failure of the partitioning method. As the number of partitions required on time increases, the task of inducing the correct global concept becomes more difficult. As the information gain available for a given partition on time is reduced, the likelihood of erroneously selecting another (possibly noisy) attribute upon which to partition the data set is increased. As a result, context changes on time are liable to be missed. Splice-2C is not affected by poor initial partitioning as it re-builds context boundaries at each iteration of contextual clustering. Hence, a poor initial partition has a minimal effect and the system can take advantage of increases in context examples. Splice-1 is still interesting, as it does substantially less work than Splice-2, and can be effective in domains with relatively few context changes. We anticipate that a stronger partitioning method would make Splice-1 more resilient to frequent changes in context. The results also indicate that the C4.5 partitioning method is not helpful in this domain. 6. Applying SPLICE to the "Learning to Fly" domain To test the Splice-2 methodology, we wished to apply it to a substantially more complex domain than the artificial data described above. We had available, data collected from flight simulation experiments used in behavioral cloning (Sammut et al., 1992). Previous work on this domain found it necessary to explicitly divide the domain into a series of individual learning tasks or stages. Splice-2 was able to induce an effective pilot for a substantial proportion of the original flight plan with no explicitly provided stages. In the following sections we briefly describe the problem domain and the application of Splice-2. 6.1. Domain The "Learning to Fly" experiments (Sammut et al., 1992) were intended to demonstrate that it is possible to build controllers for complex dynamic systems by recording the actions of a skilled operator in response to the current state of the system. A flight simulator was chosen as the dynamic system because it was a complex system requiring a high degree of skill to operate successfully and yet is well un- derstood. The experimental setup was to collect data from several human subjects flying a predetermined flight plan. These data would then be input to an induction program, C4.5. The flight plan provided to the human subjects was: 1. Take off and fly to an altitude of 2,000 feet. 2. Level out and fly to a distance of 32,000 feet from the starting point 3. Turn right to a compass heading of approximately 330 degrees. 4. At a North/South distance of 42,000 feet, turn left to head back towards the runway. The turn is considered complete when the azimuth is between 140 degrees and 180 degrees. 5. Line up on the runway. 6. Descend to the runway, keeping in line. 7. Land on the runway. The log includes 15 attributes showing position and motion, and 4 control at- tributes. The position and motion attributes record the state of the plane, whereas the control attributes record the actions of the pilot. The position and motion attributes were: on ground, g limit, wing stall, twist, elevation, azimuth, roll speed, elevation speed, azimuth speed, airspeed, climbspeed, E/W distance, altitude, N/S distance, fuel. The control attributes were: rollers, elevator, thrust and flaps. (The rudder was not used as its implementation was unrealistic.) The values for each of the control attributes provide target classes for the induction of separate decision trees for each control attribute. These decision trees are tested by compiling the trees into the autopilot code of the simulator and then "flying" the simulator. In the original experiments, three subjects flew the above flight plan times each. In all, a data set of about 90,000 records was produced. Originally, it was thought that the combined data could be submitted to the learning program. However, this proved too complex a task for the learning systems that were available. The problems were largely due to mixing data from different contexts. The first, and most critical type of context, was the pilot. Different pilots have different flying styles, so their responses to the same situation may differ. Hence, the flights were separated according to pilot. Furthermore, the actions of a particular pilot differ according to the stage of the flight. That is, the pilot adopts different strategies depending on whether he or she is turning the aircraft, climbing, landing, etc. To succeed in inducing a competent control strategy, a learning algorithm would have to be able to distinguish these different cases. Since the methods available could not do this, manual separation of the data into flight stages was required. Since the pilots were given intermediate flight goals, the division into stages was not too onerous. Not all divisions were immediately obvious. For exam- ple, initially, lining up and descending were not separated into two different stages. However, without this separation, the decision trees generated by C4.5 would miss the runway. It was not until the "line-up" stage was introduced that a successful "behavioral clone" could be produced. Until now, the stages used in behavioral cloning could only be found through human intervention which often included quite a lot of trial-and-error experimen- tation. The work described below suggests that flight stages can be treated as different contexts and that the Splice-2 approach can automate the separation of flight data into appropriate contexts for learning. 6.2. Flying with SPLICE-2 This domain introduces an additional difficulty for Splice. Previous behavioral cloning experiments have built decisions trees for each of the four actions in each of the seven stages, resulting in 28 decision trees. When flying the simulator these decision trees are switched in depending on the current stage. However, when Splice-2 is applied to the four learning tasks, viz, building a controller for elevators, another for rollers, for thrust and flaps, there is no guarantee that exactly the same context divisions will be found. This causes problems when two or more actions must be coordinated. For example, to turn the aircraft, rollers and elevators must be used together. If the contexts for these two actions do not coincide then a new roller action, say, may be commenced, but the corresponding elevator action may not start at the same time, thus causing a lack of coordination and a failure to execute the correct manoeuvre. This problem was avoided by combining rollers and elevators into a single attribute, corresponding to the stick position. Since the rollers can take one of 15 discrete values and elevators can take one of 11 discrete values, the combined attribute has 165 possible values. Of these, are represented in the training set. A further problem is how to know when to switch between contexts. The original behavioral clones included hand-crafted code to accomplish this. However, Splice builds its own contexts, so an automatic means of switching is necessary. In the on-line prediction experiment reported in Section 3.2.2, the context was selected by using a voting mechanism. This mechanism relied upon immediate feedback about classification accuracy. We do not have such feedback during a flight, so we chose learn when to switch. All examples of each stable concept were labelled with an identifier for that concept. These were input to C4.5 again, this time, to predict the context to which a state in the flight belongs, thus identifying the appropriate stable concept, which is the controller for the associated context. In designing this selection mechanism we remained with the "situation-action" paradigm that previous cloning experiments adopted so that comparisons were meaningful. We found that the original which uses only classification accuracy, did not perform well when class frequencies were wildly different. This was due to well represented classes dominating the contextual clustering process, leading to clusters with similar classification over well represented classes, and dissimilar classification over poorly represented classes. This was problematic as successful flights depend upon the correct classification of rare classes. The problem was reduced by adopting the alternative scoring method defined by Equation 2. In addition we adjusted the C4.5 parameters (pruning level) to ensure that these rare classes were not treated as noise. was also augmented to recognize domain discontinuities such as the end of one flight and the beginning of the next by altering W ij 0 such that no predictions from a flight other than the flight of example i were incorporated in any W ij 0. -45000 -40000 -35000 -30000 -25000 -20000 -15000 -10000 -5000 020004000600050015002500North/South (feet) East/West (feet) Height (feet) Figure 13. Flight comparison We have been able to successfully fly the first four stages of the flight training on data extracted from using only data from the first four stages. It should be noted that even with the changes in the domain (combining rollers and elevator) C4.5 is unable to make the first turn without the explicit division of the domain into stages. Figure 13 shows three flights: ffl The successful Splice-2 flight on stages 1 to 4. ffl The best C4.5 flight. ffl A sample complete flight. The settings used in Splice-2P were: post pruning turned off (-c 100). ffl Three iterations of the clustering stage. ffl A window size of 50 instances. ffl Initial partitioning was set to four equal divisions of the first flight. -40000 -35000 -30000 -25000 -20000 -15000 -10000 -5000 -10001000300050000100020003000 North/South (feet) East/West (feet) Height (feet) Context 3 Context 4 Figure 14. Distribution of the local concepts used in the successful flight. The chart shows only one in sixty time steps, this was necessary for clarity but does remove some brief context changes. We initially investigated the use of a larger window with a random partitioning. This successfully created two contexts: one primarily concerning the first five stages of the flight and the other concerning the last two. With this number of contexts, a successful pilot could not be induced. Reducing the window size lead to more contexts, but they were less well defined and also could not fly the plane. The solution was to bias the clustering by providing an initial partitioning based on the first four stages of the flight. Further research is needed to determine if the correspondence between the number of initial partitions and the number of flight stages is accidental or if there is something more fundamental involved. distinguished four contextual clusters with a rough correlation to flight plan stages. Each contextual cluster contains items from multiple stages of the training flights. Context 1 has instances from all four stages but has a better representation of instances from the first half of stage 1. Context 2 roughly corresponds with the second half of stage 1, stage 2, and part of 3. Instances in context 3 are from stage 2 onward, but primarily from stage 4. Context 4 is particularly interesting as it also corresponds primarily with stage 4, but contains less items than context 3 for all parts of this stage. It is not surprising that the correspondence of context to stage is not complete. The original division of the flight plan into stages was for convenience of description, and the same flight could well have been divided in many different ways. In short, the Splice contexts are capturing something additional to the division of the flight into stages. Figure 14 shows the stable concept used at each instant of the Splice-2 flight. To make context changes clearly visible, the number of points plotted in Figure 14 are fewer than were actually recorded. Because of the lower resolution, not all details of the context changes are visible in this chart. Take off was handled by context 1. Climb to 2000 feet and level out was handled by context 2. Context 3 also had occasional instances in the level out region of the flight. Context 1 again took over for straight and level flight. The right hand turn was handled by a mix of contexts 3 and 4. Subsequently flying to the North-West was handled by contexts 2 and 3 then by context 4. Initiating the left hand turn was done by context 1. The rest of the left hand turn was handled by a combination of contexts 3 and 3. When Sammut et al. (1992) divided the flight into stages, they did so by trial and error. While the partitions found were successfully used to build behavioral clones, there is no reason to believe that their partition is unique. In fact, we expect similar kinds of behaviors to be used in different stages. For example, a behavior for straight and level flight is needed in several points during the flight as is turning left or turning right. Moreover, the same control settings might be used for different maneouvres, depending on the current state of the flight. Figure 14 shows that Splice intermixed behaviors much more than was done with the manual division of stages. Although this needs further investigation, a reasonable conjecture is that an approach like Splice can find more finely tuned control strategies than can be achieved manually. In a subsequent experiment we attempted to hand craft a "perfect" switching strategy, by associating stable concepts with each stage of the flight plan. This switching method did not successfully fly the plane. At present, the addition of further stages of the flight causes catastrophic interference between the first two stages and the last 3 stages. Splice-2 is, as yet, unable to completely distinguish these parts of the flight. However, the use of Splice-2 in synthesizing controllers for stages 1 - 4 is the first time that any automated procedure has been successful for identifying contexts in this very complex domain. The use of a decision tree to select the current context is reasonably effective. As the decision tree uses only the same attributes as the stable concepts, it has no way to refer to the past. In effect, it is flying with no short term memory. This was not an issue for this work as it is a comparison with the original "Learning to Fly" project (Sammut et al., 1992) which used situation-action control. This experiment serves to demonstrate that off-line context-sensitive learning can be applied to quite complex data sets with promising results. 7. Related work There has been a substantial amount of work on dealing with known changes in context using batch learning methods. Much of this work is directly relevant to the challenges faced in using stable concepts for on-line prediction. Known context can be dealt with by dividing the domain by context, and inducing different classifiers for each context. At classification time, a meta-classifier can then be used to switch between classifiers according to the current context (Sammut et al., 1992; Katz, Gately & Collins, 1990). The application of stable concepts to on-line classification used in this paper (Sections 3.2.2 and 6.2) use a similar switching approach. Unfortunately, it is not always possible to guarantee that all hidden contexts are known. New contexts might be dealt with by adapting an existing stable concept. Kubat (1996) demonstrates that knowledge embedded in a decision tree can be transfered to a new context by augmenting the decision tree with a second tier, which is then trained on the new context. The second tier provides soft matching and weights for each leaf of the original decision tree. Use of a two tiered structure was originally proposed by Michalski (1990) for dealing with flexible concepts. Pratt (1993) shows that knowledge from an existing neural network can be re-used to significantly increase the speed of learning in a new context. These methods for the transfer of knowledge between known contexts could be used on-line to adapt stable concepts in a manner analogous to that used by Flora3. It may be possible to improve the accuracy of stable concepts by combining data from a range of contexts. Turney (1993) (Turney & Halasz, 1993) applies contextual normalization, contextual expansion and contextual weighting to a range of domains. He demonstrates that these methods can improve classification accuracy for both instance based learning (Aha, Kibler & Albert, 1991) and multivariate re- gression. This could be particularly valuable for a version of Splice using instance based methods instead of C4.5. A somewhat different on-line method designed to detect and exploit contextual attributes is MetaL(B) (Widmer, 1996). In this case, contextual attributes predict the relevance of other attributes. MetaL(B) uses any contextual attributes detected to trigger changes in the set of features presented to the classifier. While this approach and definition of context is quite different to that used by Splice, the overall philosophy is similar. Widmer concludes by stating that: ". the identification of contextual features is a first step towards naming, and thus being able to reason about, contexts." This is one of the main goals of Splice. The result of such reasoning would be a model of the hidden context. Models of hidden context could be used in on-line classification systems to augment existing reactive concept switching with a pro-active component. Models of hidden context might also be applied to improving domain understanding. Some first steps toward building models of hidden context have been taken in this article. The "Learning to Fly" experiment (Section 6.2) used a model of hidden context, based on the types of instances expected in each context, to switch between stable concepts. To summarize, Splice begins to build a bridge between on-line methods for dealing with hidden changes in context and batch methods for dealing with known change in context. 8. Conclusion This article has presented a new off-line paradigm for recognizing and dealing with hidden changes in context. Hidden changes in context can occur in any domain where the prediction task is poorly understood or where context is difficult to isolate as an attribute. Most previous work with hidden changes in context has used an on-line learning approach. The new approach, Splice , uses off-line, batch, meta-learning to extract hidden context and induce the associated stable concepts. It incorporates existing machine learning systems (in this paper, C4.5 (Quinlan, 1993)). The initial implementation, was briefly reviewed and a new version, Splice-2, presented in full. The evaluation of the Splice approach included an on-line prediction task, a series of hidden context recognition tasks, and a complex control task. was the initial implementation of the Splice approach and used C4.5 to divide a data series by likely changes of context. A process called contextual clustering then grouped intervals appearing to be from the same context. This process used the semantics of concepts induced from each interval as a measure of the similarity of context. The resulting contextual clusters were used to create context specific concepts and to specify context boundaries. limitations of Splice-1 by permitting refinement of partition boundaries. Splice-2 clusters on the basis of individual members of the data series. Hence, context boundaries are not restricted to the boundaries found in the partitioning stage and context boundaries can be refined. Splice-2 is much more robust to the quality of the initial partitioning. successfully detected and dealt with hidden context in a complex control task. "Learning to Fly" is a behavioral cloning domain based upon learning an autopilot given a series of sample flights with a fixed flight plan. Previous work on this domain required the user to specify stages of the flight. Splice-2 was able to successfully fly a substantial fragment of the initial flight plan without these stages (or contexts) being specified. This is the first time that any automated procedure has been successful for identifying context in this very complex domain. A number of improvements could be made to the Splice algorithms. The partitioning method used was shown to be problematic for Splice-1 at high levels of noise and hidden changes in context. While the use of an existing machine learning system to provide partitioning is elegant, a better solution may be to implement a specialized method designed to deal with additional complexity over time. One approach to this is to augment a decision tree algorithm to allow many splits (Fayyad Irani, 1993) on selected attributes. Neither Splice-1 nor Splice-2 provide a direct comparison of the relative advantage of dividing the domain into one set of contexts over another. One comparison method that could be used is the minimum description length (MDL) heuristic (Rissanen, 1983). The MDL principle is that the best theory for a given concept will minimize the amount of information that need be sent from a sender to a receiver so that the receiver can correctly classify items in a shared dataset. In this case, the information to be sent must contain any stable concepts, a context switching method and a list of exceptions. At the very least, this would allow the direct comparison of a given context-sensitive global concept (using stable concepts and context switching) with a context-insensitive global concept. Further, a contextual clustering method could use an MDL heuristic to guide a search through the possible context divisions. The approaches used here for selecting the current context were an on-line voting method for domains with immediate feedback and a decision tree for a domain without immediate feedback. More sophisticated approaches would use a model of the hidden context. Such a model could use knowledge about the expected context duration, order and stability. It could also incorporate other existing attributes and domain feedback. The decision tree used for context switching in the learning to fly task is a primitive implementation of such a model using only existing attributes to select the context. An exciting possibility is to use the characteristics of contexts identified by Splice to guide a search of the external world for an attribute with similar characteristics. Any such attributes could then be incorporated with the current attribute set allowing a bootstrapping of the domain representation. This could be used within the Knowledge Discovery in Databases (KDD) approach (Fayyad, Piatsky-Shapiro Smyth, 1996) which includes the notion that analysts can reiterate the data selection and learning (data mining) tasks. Perhaps too, this method could provide a way for an automated agent to select potentially useful attributes from the outside world, with which to extend its existing domain knowledge. Acknowledgments We would like to thank the editors and the anonymous reviewers. Their suggestions led to a much improved paper. Michael Harries was partially supported by an Australian Postgraduate Award (Industrial) sponsored by RMB Australia. Notes 1. In the following experiments, n% noise implies that the class was randomly selected with a probability of n%. This method for generating noise was chosen to be consistent with Widmer and Kubat (1996). --R Incremental batch learning. Detecting concept drift in financial time series prediction using symbolic machine learning. Neural Computation Floating approximation in time-varying knowledge bases Second tier for decision trees. Adapting to drift in continuous domains. Classifying sensor data with CALCHAS. Learning flexible concepts: Fundimental ideas and a method based on two-tiered representation Some experiments in applying inductive inference principles to surface reconstruction. A universal prior for integers and estimation by minimum description length. Annals of Statistics Density adaptive learning and forgetting. Learning to fly. Robust classification with context-sensitive features Recognition and exploitation of contextual clues via incremental meta- learning Effective learning in dynamic environments by explicit concept tracking. Learning in the presence of concept drift and hidden contexts. --TR --CTR Francisco Ferrer-Troyano , Jesus S. Aguilar-Ruiz , Jose C. Riquelme, Incremental rule learning based on example nearness from numerical data streams, Proceedings of the 2005 ACM symposium on Applied computing, March 13-17, 2005, Santa Fe, New Mexico Francisco Ferrer-Troyano , Jesus S. Aguilar-Ruiz , Jose C. Riquelme, Data streams classification by incremental rule learning with parameterized generalization, Proceedings of the 2006 ACM symposium on Applied computing, April 23-27, 2006, Dijon, France Anand Narasimhamurthy , Ludmila I. Kuncheva, A framework for generating data to simulate changing environments, Proceedings of the 25th conference on Proceedings of the 25th IASTED International Multi-Conference: artificial intelligence and applications, p.384-389, February 12-14, 2007, Innsbruck, Austria Chunsheng Yang , Sylvain Ltourneau, Learning to predict train wheel failures, Proceeding of the eleventh ACM SIGKDD international conference on Knowledge discovery in data mining, August 21-24, 2005, Chicago, Illinois, USA Dwi H. Widyantoro , John Yen, Relevant Data Expansion for Learning Concept Drift from Sparsely Labeled Data, IEEE Transactions on Knowledge and Data Engineering, v.17 n.3, p.401-412, March 2005 Antonin Rozsypal , Miroslav Kubat, Association mining in time-varying domains, Intelligent Data Analysis, v.9 n.3, p.273-288, May 2005 Bruce Edmonds, Learning and exploiting context in agents, Proceedings of the first international joint conference on Autonomous agents and multiagent systems: part 3, July 15-19, 2002, Bologna, Italy

concept drift;hidden context;contextual clustering;batch learning;context-sensitive learning

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Tracking the Best Disjunction.

Littlestone developed a simple deterministic on-line learning algorithm for learning k-literal disjunctions. This algorithm (called {\sc Winnow}) keeps one weight for each of then variables and does multiplicative updates to its weights. We develop a randomized version of {\sc Winnow} and prove bounds for an adaptation of the algorithm for the case when the disjunction may change over time. In this case a possible target {\it disjunction schedule} &Tgr; is a sequence of disjunctions (one per trial) and the {\it shift size} is the total number of literals that are added/removed from the disjunctions as one progresses through the sequence.We develop an algorithm that predicts nearly as well as the best disjunction schedule for an arbitrary sequence of examples. This algorithm that allows us to track the predictions of the best disjunction is hardly more complex than the original version. However, the amortized analysis needed for obtaining worst-case mistake bounds requires new techniques. In some cases our lower bounds show that the upper bounds of our algorithm have the right constant in front of the leading term in the mistake bound and almost the right constant in front of the second leading term. Computer experiments support our theoretical findings.

Introduction One of the most significant successes of the Computational Learning Theory community has been Littlestone's formalization of an on-line model of learning and the development of his algorithm Winnow for learning disjunctions (Littlestone, 1989, 1988). The key feature of Winnow is that when learning disjunctions of constant size, the number of mistakes of the algorithm grows only logarithmically with the input dimension. For many other standard algorithms such as the Perceptron Algorithm (Rosenblatt, 1958), the number of mistakes can grow linearly in the dimension (Kivinen, Warmuth, & Auer, 1997). In the meantime a number of algorithms similar to Winnow have been developed that also show the logarithmic growth of the loss bounds in the dimension (Littlestone & Warmuth, 1994; Vovk, 1990; Cesa-Bianchi et al., 1997; Haussler, Kivinen, & Warmuth, 1994). * An extended abstract appeared in (Auer & Warmuth, 1995). * M. K. Warmuth acknowledges the support of the NSF grant CCR 9700201. In this paper we give a refined analysis of Winnow, develop a randomized version of the algorithm, give lower bounds that show that both the deterministic and the randomized version are close to optimal, and adapt both versions so that they can be used to track the predictions of the best disjunction. Consider the following by now standard on-line learning model (Littlestone, 1989, 1988; Vovk, 1990; Cesa-Bianchi et al., 1997). Learning proceeds in trials. In trial the algorithm is presented with an instance x t (in our case an n-dimensional binary vector) that is used to produce a binary prediction - y t . The algorithm then receives a binary classification y t of the instance and incurs a mistake if - The goal is to minimize the number of mistakes of the algorithm for an arbitrary sequence of examples h(x This is of course a hopeless scenario: for any deterministic algorithm an adversary can always choose the sequence so that the algorithm makes a mistake in each trial. A more reasonable goal is to minimize the number of mistakes of the algorithm compared to the minimum number of mistakes made by any concept from a comparison class. 1.1. The (non-shifting) basic setup In this paper we use monotone 1 k-literal disjunctions as the comparison class. If the dimension (number of Boolean attributes/literals) is n then such disjunctions are Boolean formulas of the form x , where the (distinct) indices i j lie in ng. The number of classification errors of such a disjunction with respect to a sequence of examples is simply the total number of misclassifications that this disjunction produces on the sequence. The goal is to develop algorithms whose number of mistakes is not much larger than the number of classification errors of the best disjunction, for any sequence of examples. In this paper we consider the case where the mistakes of the best ("target") disjunction are caused by attribute errors. The number of attribute errors of an example with respect to a target disjunction u is the minimum number of attributes/bits of x that have to be changed so that for the resulting x 0 , y. The number of attribute errors for a sequence of examples with respect to a target concept is simply the total number of such errors for all examples of the sequence. Note that if the target u is a k-literal monotone disjunction then the number of attribute errors is at most k times the number of classification errors with respect to u (i.e. k times the number of examples (x; y) in the sequence for which u(x) 6= y). Winnow can be tuned as a function of k so that it makes at most O(A+k ln(n=k)) mistakes on any sequence of examples where the best disjunction incurs at most A attribute errors (Littlestone, 1988). We give a randomized version of Winnow and give improved tunings of the original algorithm. The new algorithm can be tuned based on k and A so that its expected mistake bound is at most A any sequence of examples for which there is a monotone k-literal disjunction with at most A attribute errors. We also show how the original deterministic algorithm can be tuned so that its number of mistakes is at most 2A for the same set of sequences. Our lower bounds show that these bounds are very close to optimal. We show that for any algorithm the expected number of mistakes must be at least our upper bound has the correct constant on the leading term and almost the optimal constant on the second term. For deterministic algorithms our lower bounds show that the constant on the leading term is optimal. Our lower bounds for both the deterministic and the randomized case cannot be improved significantly because there are essentially matching upper bounds achieved by non-efficient algorithms with the correct factors on the first and the second term. These algorithms use experts (Cesa-Bianchi et al., 1997): each expert simply computes the value of a particular k-literal disjunction and one weight is kept per expert. This amounts to expanding the n-dimensional Boolean inputs into Boolean inputs and then using single literals (=experts) (Littlestone & Warmuth, 1994; Vovk, 1990; Cesa-Bianchi et al., 1997) as the comparison class instead of k- literal disjunctions. The expected number of mistakes of the randomized algorithm is at most Q+ a bound on the number of classification errors of the best k-literal disjunction. The mistake bound of the deterministic algorithm is exactly twice as high. Observe that these algorithms have to use about n k weights, and that they need that much time in each trial to calculate their prediction and update the weights. Thus their run time is exponential in k. In contrast, our algorithm uses only n weights. On the other hand the noise in the upper bounds of our efficient algorithm is measured in attribute errors rather than classification errors. This arises since we are using just one weight per attribute. Recall that a classification error with respect to a k-literal disjunction can equate to up to k attribute errors. To capture errors that affect up to k attributes efficiently the expansion to experts seems to be unavoidable. Nevertheless, it is surprising that our version of Winnow is able to get the right factor before the number of attribute errors A and for the randomized version almost the right factor before the square root term. In some sense Winnow compresses \Delta weights to only n weights. At this point we don't have a combinatorial interpretation of our weights. Such an interpretation was only found for the single literal (expert) case (Cesa-Bianchi, Freund, Helmbold, & Warmuth, 1996). As Littlestone (Littlestone, 1991) we use an amortized analysis with an entropic potential function to obtain our worst-case loss bounds. However besides the more careful tuning of the bounds we take the amortized analysis method a significant step further by proving mistake bounds of our algorithm as compared to the best shifting disjunction. 1.2. Shifting disjunctions Assume that a disjunction u is specified by a n-dimensional binary vector, where the components with value 1 correspond to the monotone literals of the disjunction. For two disjunctions u and u 0 the Hamming distance measures how many literals have to be "shifted" to obtain u 0 from u. A disjunction schedule T for a sequence of examples of length T is simply a sequence of T disjunctions u t . The (shift) size of the schedule T is is the all zero vector). In the original non-shifting case all u are equal to some k-literal disjunction u and accordingly to the above definition the "shift size" is k. At trial t the schedule T predicts with disjunction u t . We define the number of attribute errors of an example sequence h(x t ; y t )i with respect to a schedule T as the total number of attributes that have to be changed in the sequence of examples to make it consistent with the schedule T , i.e. for which the changed instances x 0 Note that the loss bounds for the non-shifting case can be written as cA+O( \Delta is the number of bits it takes to describe a disjunction with k literals, and where for the randomized and for the deterministic algorithm. Surprisingly, we were able to prove bounds of the same form for the shifting disjunction case. B is now the number of bits it takes to describe the best schedule T and A is the number of attribute errors of this schedule. If Z is the shift size of schedule T then it takes log 2 Z bits to describe a schedule T in respect to a given sequence of examples. 2 Our worst-case mistake bounds are similar to bounds obtained for "competitive algorithms" in that we compare the number of mistakes of our algorithm against the number of attribute errors of the best off-line algorithm that is given the whole sequence ahead of time. The off-line algorithm still incurs A attribute errors and here we bound the additional loss of the on-line algorithm over the number of attribute errors of the best schedule (as opposed to the coarser method of bounding the ratio of on-line over off-line). Winnow does multiplicative updates to its weights. Whenever the algorithm makes a mistake then the weights of all the literals for which the corresponding bit in the current input instance is one are multiplied by a factor. In the case of Winnow2, the version of Winnow this paper is based on (Littlestone, 1988), this factor is either ff or 1=ff, where ff ? 1 is a parameter of the algorithm. The multiplicative weight updates might cause the weights of the algorithm to decay rather rapidly. Since any literal might become part of the disjunction schedule even when it was misleading during the early part of the sequence of examples, any algorithm that is to predict well as compared to the best disjunction schedule must be able to recover weights quickly. Our extension of Winnow2 simply adds a step to the original algorithm that resets a weight to fi=n whenever it drops below this boundary. Similar methods for lower bounding the weights were used in the algorithm Wml of (Littlestone & Warmuth, 1994) which was designed for predicting as well as the best shifting single literal (which is called expert in (Cesa-Bianchi et al., 1997)). In addition to generalizing the work of (Littlestone & Warmuth, 1994) to arbitrary size disjunctions we were able to optimize the constant in the leading term of the mistake bound of Winnow and develop a randomized version of the algorithm. In (Herbster & Warmuth, 1998) the work of (Littlestone & Warmuth, 1994) was generalized in a different direction. The focus there is to predict as well as the best shifting expert, where "well" is measured in terms of other loss functions than the discrete loss (counting mistakes) which is the loss function used in this paper. Again the basic building block is a simple on-line algorithm that uses multiplicative weight updates (Vovk, 1990; Haussler et al., 1994) but now the predictions and the feedback in each trial are real-valued and lie in the interval [0; 1]. The class of loss functions includes the natural loss functions of log loss, square loss and Hellinger loss. Now the loss does not occur in "large" discrete units. Instead the loss in a trial my be arbitrarily small and thus more sophisticated methods are needed for recovering small weights quickly (Herbster & Warmuth, 1998) than simply lower bounding the weights. Why are disjunctions so important? Whenever a richer class is built by (small) unions of a large number of simple basic concepts, our methods can be applied. Simply expand the original input into as many inputs as there are basic concepts. Since our mistake bounds only depend logarithmically on the number of basic con- cepts, we can even allow exponentially many basic concepts and still have polynomial mistake bounds. This method was previously used for developing noise robust algorithms for predicting nearly as well as the best discretized d-dimensional axis-parallel box (Maass & Warmuth, 1998; Auer, 1993) or as well as the best pruning of a decision tree (Helmbold & Schapire, 1997). In these cases a multiplicative algorithm maintains one weight for each of the exponentially many basic concepts. However for the above examples, the multiplicative algorithms with the exponentially many weights can still be simulated efficiently. Now, for example, the methods of this paper immediately lead to an efficient algorithm for predicting as well as the best shifting d-dimensional box. Thus by combining our methods with existing al- gorithms, we can design efficient learning algorithms with provably good worst-case loss bounds for more general shifting concepts than disjunctions. Besides doing experiments on practical data that exemplify the merits of our worst-case mistake bounds, this research also leaves a number of theoretical open prob- lems. Winnow is an algorithm for learning arbitrary linear threshold functions and our methods for tracking the best disjunction still need to be generalized to learning this more general class of concepts. We believe that the techniques developed here for learning how to predict as well as the best shifting disjunction will be useful in other settings such as developing algorithms that predict nearly as well as the best shifting linear combination. Now the discrete loss has to be replaced by a continuous loss function such as the square loss, which makes this problem more challenging. 1.3. Related work There is a natural competitor to Winnow which is the well known Perceptron algorithm (Rosenblatt, 1958) for learning linear threshold functions. This algorithm does additive instead of multiplicative updates. The classical Perceptron Convergence Theorem gives a mistake bound for this algorithm (Duda & Hart, 1973; Haykin, 1994), but this bound is linear in the number of attributes (Kivinen et al., 1997) whereas the bounds for the Winnow-like algorithms are logarithmic in the number of attributes. The proof of the Perceptron Convergence Theorem can also be seen as an amortized analysis. However the potential function needed for the perceptron algorithm is quite different from the potential function used for the analysis of Winnow. If w t is the weight vector of the algorithm in trial t and u is a target weight vector, then for the perceptron algorithm 2 is the potential function where jj:jj 2 is the Euclidean length of a vector. In contrast the potential function used for the analysis of Winnow (Littlestone, 1988, 1989) that is also used in this paper is the following generalization 3 of relative entropy (Cover, In the case of linear regression a framework was developed (Kivinen & Warmuth, 1997) for deriving updates from the potential function used in the amortized anal- ysis. The same framework can be adapted to derive both the Perceptron algorithm and Winnow. The different potential functions for the algorithms lead to the additive and multiplicative algorithms, respectively. The Perceptron algorithm is seeking a weight vector that is consistent with the examples but otherwise minimizes some Euclidean length. Winnow instead minimizes a relative entropy and is thus rooted in the Minimum Relative Entropy Principle of Kullback (Kapur & Kesavan, 1992; Jumarie, 1990). 1.4. Organization of the paper In the next section we formally define the notation we will use throughout the paper. Most of them have already been discussed in the introduction. Section 3 presents our algorithm and Section 4 gives the theoretical results for this algorithm. In Section 5 we consider some more practical aspects, namely how the parameters of the algorithm can be tuned to achieve good performance. Section 6 reports some experimental results. The analysis of our algorithm and the proofs for Section 4 are given in Section 7. Lower bounds on the number of mistakes made by any algorithm are shown in Section 8 and we conclude in Section 9. 2. Notation A target schedule a sequence of disjunctions represented by n-ary bit vectors u . The size of the shift from disjunction u to disjunction u t is z j. The total shift size of schedule T is z t where we assume that u To get more precise bounds for the case when there are shifts in the target schedule we will distinguish between shifts where a literal is added to the disjunction and shifts where a literal is removed from the disjunction. Thus we define z t as the number of times a literal is switched on, and t as the number of times a literal is switched off. A sequence of examples consists of attribute vectors classifications y t 2 f0; 1g. The prediction of disjunction u t for attribute vector x t is u t The number of attribute errors a t at trial t with respect to a target schedule T is the minimal number of attributes that have to be changed, resulting in x 0 t , such that u t That is a g. The total number of attribute errors of sequence S with respect to schedule T is a t . We denote by S(Z; A; n) the class of example sequences S with n attributes which are consistent with some target schedule T with shift size Z and with at most A attribute errors. If we wish to distinguish between positive and negative shifts we denote the corresponding class by S(Z are the numbers of literals added and removed, respectively, in the target schedule. By S 0 (k; A; n) we denote the class of example sequences S with n attributes which are consistent with some non-shifting target schedule (i.e. and with at most A attribute errors. For the case that only upper bounds on Z, Z are known we denote the corresponding classes by S - (Z; A; z-Z -k The loss of a learning algorithm L on an example sequence S is the number of misclassifications is the binary prediction of the learning algorithm L in trial t. 3. The algorithm We present algorithm Swin ("Shifting Winnow"), see Table 1, an extension of Littlestone's Winnow2 algorithm (Littlestone, 1991). Our extension incorporates a randomization of the algorithm, and it guarantees a lower bound on the weights used by the algorithm. The algorithm maintains a vector of n weights for the n attributes. By w we denote the weights at the end of trial t, Table 1. Algorithm Swin Parameters: The algorithm uses parameters ff ? Initialization: Set the weights to initial values In each trial t - 1 set r predict ae Receive the binary classification y t . If y Update: If y t 6= p(r t ) then for all 1. w 0 2. w and w 0 denotes the initial value of the weight vector. In trial t the algorithm predicts using the weight vector w . The prediction of the algorithm depends on r 1]. The algorithm predicts 1 with probability p(r t ), and it predicts 0 with probability obtain a deterministic algorithm one has to choose a function p predicting the algorithm receives the classification y t . If y i.e. the weight vector is not modified. Since y t 2 f0; 1g and p(r t ) 2 [0; 1] this can only occur when the prediction was deterministic, i.e. p(r t correct. An update occurs in all other cases when the prediction was wrong or p(r t The updates of the weights are performed in two steps. The first step is the original Winnow update, and the second step guarantees that no weight is smaller than fi for some parameter fi (a similar approach was taken in (Littlestone & Warmuth, 1994)). Observe that the weights are changed only if the probability of making a mistake was non-zero. For the deterministic algorithm this means that the weights are changed only if the algorithm made a mistake. Furthermore the i-th weight is modified only if x 1. The weight is increased (multiplied by ff) if y it is decreased (divided by ff) if y The parameters ff, fi, w 0 , and the function p(\Delta), have to be set appropriately. A good choice for function p(\Delta) is the following: for a randomized prediction let if (RAND) and for a deterministic version of the algorithm let (DET) For the randomized version one has to choose fi ! ln ff . Observe that (DET) is obtained from (RAND) by choosing the threshold in (RAND). This corresponds to the straightforward conversion from a randomized prediction algorithm into a deterministic prediction algorithm. Theoretically good choices of the parameters ff, fi, and w 0 are given in the next section and practical issues for tuning the parameters are discussed in Section 5. 4. Results In this section we give rigorous bounds on the (expected) number of mistakes of Swin, first in general and then for specific choices of ff, fi, and w 0 , all with p(\Delta) chosen from (RAND) or (DET). These bounds can be shown to be close to optimal for adversarial example sequences, for details see Section 8. Theorem 1 (randomized version) Let ff ? 1, n , and p(\Delta) as in (RAND). Then for all S A If fi - n e then the bound holds for all S Theorem 2 (deterministic version) Let ff ? 1, n , and p(\Delta) as in (DET). Then for all S A If fi - n e then the bound holds for all S Theorem 3 (non-shifting case) Let ff ? 1, if Swin uses the function p(\Delta) given by (RAND), and if Swin uses the function p(\Delta) given by (DET). e then the bounds hold for all S 2 S - Remark. The usual conversion of a bound M for the randomized algorithm into a bound for the deterministic algorithm would give 2M as the deterministic bound. 4 But observe that our deterministic bound is just 1 times the randomized bound. Since at any time a disjunction cannot contain more than n literals we have Z which gives the following corollary. , and w . If p(\Delta) as in (RAND) then for all S 2 S - (Z; A; n) A j. j. If p(\Delta) as in (DET) then for all S 2 S - (Z; A; n) A ffn j. j. At first we give results on the number of mistakes of Swin, if no information besides n, the total number of attributes, is given. Theorem 4 Let n , and p(\Delta) be as in (RAND). Then for all S 2 S - (Z; A; n) n , and p(\Delta) be as in (DET). Then for all S 2 n , and p(\Delta) be as in (RAND). Then for all S 2 then the above bound holds for all S 2 S - n). For n - 2 we have , and p(\Delta) be as in (DET). Then for all S 2 S 0 (k; A; n) then the above bound holds for all S 2 S - n). For n - 2 we have In Section 8 we will show that these bounds are optimal up to constants. If A and Z are known in advance then the parameters of the algorithm can be tuned to obtain even better results. If for example in the non-shifting case the number k of attributes in the target concept is known we get Theorem 5 Let n , and p(\Delta) be as in (RAND). Then for e then the above bound holds for all S 2 S - n). For k - n e we set e and get EM(Swin;S) - 1:44 e for all S 2 S - , and p(\Delta) be as in (DET). Then for all S 2 S 0 (k; A; n) e then the above bound holds for all S 2 S - n). For k - n e we set e and get M(Swin;S) - 2:75 e for all S 2 S - Of particular interest is the case when A is the dominant term, i.e. A AE k ln n Theorem 6 Let A - k ln n A n , and p(\Delta) be as in (RAND). Then for all S 2 S 0 (k; A; n) r e then the above bound holds for all S 2 S - n). For k - n e , A - n e , e , we have EM(Swin;S) - A e for all If A - 2k A , and p(\Delta) be as in (DET), then for all S 2 S 0 (k; A; n) r e then the above bound holds for all S 2 S - n). For k - n e e e , we have M(Swin;S) - 2A e for all In the shifting case we get for dominant A AE Z ln n Theorem 7 Let Z+minfn;Zg Z , and A and Zsuch that ffl - 1 . Then for n , and p(\Delta) as in (RAND), and for all S 2 S - (Z; A; n), r An Z Z+minfn;Zg A An Z n , and p(\Delta) as in (DET), then for all S 2 S - (Z; A; n), r An Z In Section 8 we will show that in Theorems 6 and 7 the constants on A are optimal. Furthermore we can show for the randomized algorithm that also the magnitude of the second order term in Theorem 6 is optimal. 5. Practical tuning of the algorithm In this section we give some thoughts on how the parameters ff, fi, and w 0 of Swin should be chosen for particular target schedules and sequences of examples. Our recommendations are based on our mistake bounds which hold for any target schedule and for any sequence of examples with appropriate bounds on the number of shifts and attribute errors. Thus it has to be mentioned that, since many target schedules and many example sequences are not worst case, our bounds usually overestimate the number of mistakes made by Swin. Therefore parameter settings different from our recommendations might result in a smaller number of errors for a specific target schedule and example sequence. On the other hand Swin is quite insensitive to small changes in the parameters (see Section 6) and the effect of such changes should be benign. If little is known about the target schedule or the example sequence than the parameter settings of Theorems 4 or 5 are advisable since they balance well between the effect of target shifts and attribute errors. If good estimates for the number of target shifts and the number of attribute errors are known than good parameters can be calculated by numerically minimizing the bounds in Theorems 1, 2, 3 or Corollary 1, respectively. If the average rate of target shifts and attribute errors is known such that Z - r Z T and A - r AT with r Z ? 0; r A - 0 then for large T the error rate r M(Swin;S)=T is by Corollary 1 approximately upper bounded by r A for randomized predictions and by r A ffn for deterministic predictions. Again, optimal choices for ff and fi can be obtained by numerical minimization. 6. Experimental results The experiment reported in this section is not meant to give a rigorous empirical evaluation of algorithm Swin. Instead, it is intended as an illustration of the typical behavior of Swin, compared with the theoretical bound and also with a version of Winnow which was not modified to adapt to shifts in the target schedule. In our experiment we used attributes and a target schedule T of length which starts with 4 active literals. After 1000 trials one of the literals is switched off and after another 1000 trials another literal is switched on. This switching on and switching off of literals continues as depicted in Figure 1. Thus there are initially 4 active literals). The example sequence h(x t ; y t )i was chosen such that for half of the examples y and for the other half y The values of attributes not appearing in the target schedule were chosen at random such that x probability 1/2. For examples with y exactly one of the active attributes (chosen at random) was set to number of trials number of active literals Figure 1. Shifts in the target schedule used in the experiment. 1. For examples with attribute errors all relevant attributes were either set to 1 (for the case y set to 0 (for the case y Attribute errors occurred at trials with y 1 and at trials a Figure 2 shows the performance of Swin compared with the theoretical bound where the parameters were set by numerically minimizing the bound of Corollary 1 as described in the previous section, which yielded The theoretical bound at trial t is calculated from the actual number of shifts and attribute errors up to this trial. Thus an increase of the bound is due to a shift in the target schedule or an attribute error at this trial. In Figure 2 the reasons for these increases are indicated by z + for a literal switched on, z \Gamma for a literal switched off, and a for attribute errors. Figure 2 shows that the theoretical bound very accurately depicts the behavior of Swin, although it overestimates the actual number of mistakes by some amount. It can be seen that switching off a literal causes far less mistakes than switching on a literal, as predicted by the bound. Also the relation between attribute errors and mistakes can be seen. The performance of Swin for the whole sequence of examples is shown in Figure 3 and it is compared with the performance of a version of Winnow which was not modified for target shifts. As can be seen Swin adapts very quickly to shifts in the target schedule. On the other hand, the unmodified version of Winnow makes more and more mistakes for each shift. The unmodified version of Winnow we used is just Swin with Thus the weights are not lower bounded and can become arbitrarily small which causes a large number of mistakes if the corresponding literal becomes active. We used the z - a a=4 z number of trials number of mistakes theoretical bound performance of SWIN Figure 2. Comparison of Swin with the theoretical bound for a particular target schedule and sequence of examples. Shifts and attribute errors are indicated by z number of trials number of mistakes theoretical bound performance of SWIN performance of Winnow Figure 3. A version of Winnow which does not lower bound the weights makes many more mistakes than Swin. same ff for the unmodified version but we set w which is optimal for the initial part of the target schedule. Therefore the unmodified Winnow adapts very quickly to this initial part, but then it makes an increasing number of mistakes for each shift in the target schedule. For each shift the number of mistakes made approximately doubles. In the last plot, Figure 4, we compare the performance of Swin with tuned parameters to the performance of Swin with the generic parameter setting given by Theorem 4. Although the tuned parameters do perform better the difference is number of trials number of mistakes theoretical bound performance of SWIN Figure 4. Tuned parameters of Swin versus the generic parameters relatively small. The overall conclusion of our experiment is that, first, the theoretical bounds capture the actual performance of the algorithm quite well, second, that some mechanism of lower bounding the weights of Winnow is necessary to make the algorithm adaptive to target shifts, and third, that moderate changes in the parameters do not change the qualitative behavior of the algorithm. 7. Amortized analysis In this section we first prove the general bounds given in Theorems 1 and 2 for the randomized and for the deterministic version of Swin. Then from these bounds we calculate the bounds given in Theorems 4-7 for specific choices of the parameters. The analysis of the algorithm proceeds by showing that the distance between the weight vector of the algorithm w t , and vector u t representing the disjunction at trial t, decreases, if the algorithm makes a mistake. The potential/distance function used for the previous analysis of Winnow (Littlestone, 1988, 1989, 1991) is the following generalization of relative entropy to arbitrary non-negative weight vectors: (This distance function was also used for the analysis of the Egu regression algorithm (Kivinen & Warmuth, 1997), which shows that Winnow is related to the algorithm.) By taking derivatives it is easy to see that the distance is minimal and equal to 0 if and only if w . With the convention that 0 and the assumption that u 2 f0; 1g n the distance function simplifies to We start with the analysis of the randomized algorithm with shifting target dis- junctions. The other cases will be derived easily from this analysis. At first we calculate how much the distance D(u t ; w t ) changes between trials: Observe that term (1) might be non-zero in any trial, but that terms (2) and (3) are non-zero only if the weights are updated in trial t. For any fl, ffi with can lower bound term (1) by If the weights are updated in trial t, term (2) is bounded by The third equality holds since each x t;i 2 f0; 1g. Remember that x 0 t is obtained from x t by removing the attribute errors from x t . The last inequality follows from the fact that At last observe that w t;i 6= w 0 only if y . In this case w 0 ffn and we get for term (3) Summing over all trials we have to consider the trials where the weights are updated and we have to distinguish between trials with y denote these trials. Then by the above bounds on terms (1), (2), and (3) we have \Theta r t Now we want to lower bound the sum over M 0 and M 1 by the expected (or total) number of mistakes of the algorithm. We can do this by choosing an appropriate function p(\Delta; w t ). We denote by p t the probability that the algorithm makes a mistake in trial t. Then the expected number of mistakes is . Observe that since in this case y Thus it is sufficient to find a function p(\Delta) and a constant C with and For such a function p(\Delta) satisfying (4) and (5) we get assuming that S 2 S(Z we can upper bound the expected number of mistakes by Hence we want to choose p(\Delta) and C such that C is as big as possible. For that fix p(\Delta) and let r be a value where p(r ) becomes 1. 5 Since the left hand sides of equations (4) and (5) are continuous we get (r combining these two we have ff This can be achieved by choosing p(\Delta) as in (RAND) which satisfies (4) and (5) for course we have to choose fi ! ln ff . Putting everything together we have the following lemma. and assume that are the weights used by algorithm Swin. Then for all S if Swin uses the function p(\Delta) given by (RAND). For the deterministic variant of our algorithm the function p(\Delta) has to take values in f0; 1g. Thus we get from (4) and (5) that (r\Gammafi)(1\Gamma1=ff) - C and r(1\Gammaff)+ln ff - C which yields This we get by choosing p(\Delta) as in (DET) which satisfies (4) and (5) for and assume that are the weights used by algorithm Swin. Then for all S if Swin uses the function p(\Delta) given by (DET). Now we are going to calculate bounds fl; ffi on 1 We get these bounds by lower and upper bounding w t;i . Obviously w t;i - fi for all t and i. The upper bound on w t;i is derived from the observation that w t;i ? w t\Gamma1;i only if y with the p(\Delta) as in (RAND) or (DET), and r t - w t\Gamma1;i x t;i we find that w t;i - ff. Thus ln efi Lemma 3 If fi the weights w t;i of algorithm Swin with function p(\Delta) as in (RAND) or (DET) satisfy and Proof of Theorems 1 and 2. By Lemmas 1, 2, and 3. 2 Proof of Theorem 3 In the non-shifting case where u and it is is the number of attributes in the target disjunction u. Thus in the non-shifting case the term in the upper bounds of Lemmas 1 and 2 can be replaced by k , which gives the theorem. 2 7.1. Proofs for specific choices of the parameters Proofs of Theorems 4 and 5. By Theorem 3 and Corollary 1 and simple calcu- lations. 2 Proof of Theorem 6. For we get from Theorem 3 that with with 2. Then In the second inequality we used that ffl - 1. Substituting the values for c and ffl gives the statements of the theorem. 2 Proof of Theorem 7. For we get from Corollary 1 that j. j. with j. j. with 2. Then for ffl - 1=10 j. j. c c for r A An Z This gives the bounds of the theorem. 2 8. Lower bounds We start by proving a lower bound for the shifting case. We show that for any learning algorithm L there are example sequences S for which the learning algorithm makes "many" mistakes. Although not expressed explicitly in the following theorems we will show that these sequences S can be generated by target schedules each disjunction u t consists of exactly one literal, i.e. is the jth unit vector. Our first lower bound shows that for any deterministic algorithm there is an adversarial example sequence in S(Z; A; n) such that it makes at least 2A many mistakes. Related upper bounds are given in Theorems 4 and 7. Theorem 8 For any deterministic learning algorithm L, any n - 2, any Z - 1, and any A - 0, there is an example sequence S 2 S(Z; A; n) such that Proof. For notational convenience we assume that R - 1. We construct the example sequence S depending on the predictions of the learning algorithm such that the learning algorithm makes a mistake in each trial. We partition the trials into R rounds. The first R \Gamma 1 rounds have length -, the last round has length - errors will occur only within the last trials. We choose the target schedule such that during each round the target disjunction does not change and is equal to some e j . At the beginning of each round there are disjunctions consistent with the examples of this round. After l trials in this round there are still 2 - \Gammal consistent disjunctions: we construct the attribute vector by setting half of the attributes which correspond to consistent disjunctions to 1, and the other attributes to 0. Furthermore we set y y t is the prediction of the algorithm for this attribute vector. Obviously half of the disjunctions are consistent with this example, and thus the number of consistent disjunctions is divided by 2 in each trial. Thus in each of the first R \Gamma 1 rounds there is a disjunction consistent with all - examples of this round. After in the last round there are two disjunctions consistent with the examples of this round. For the remaining 2A+1 trials we fix some attribute vector for which these two disjunctions predict differently, and again we set y y t . Thus one of these disjunctions disagrees at most A times with the classifications in these This disagreement can be seen as caused by A attribute errors, so that the disjunction is consistent with all the examples in the last round up to A attribute errors. 2 Remark. Observe that a lower bound for deterministic algorithms like implies the following lower bound on randomized algorithms: This follows from the fact that any randomized learning algorithm can be turned into a deterministic learning algorithm which makes at most twice as many mistakes as the randomized algorithm makes in the average. This means that Theorem 8 implies for any randomized algorithm L that there are sequences S 2 S(Z; A; n) with Remark. As an open problem it remains to show lower bounds that have the same form as the upper bounds of Theorem 7 with the square root term. Now we turn to the non-shifting case. For are already lower bounds known. Lemma 4 (Littlestone & Warmuth, 1994) For any deterministic learning algorithm L, any n - 2, and any A - 0, there is an example sequence S 2 S 0 (1; A; n) such that A slight modification of results in (Cesa-Bianchi et al., 1997) gives Lemma 5 (Cesa-Bianchi et al., 1997) There are functions n(j) and A(n; j) such that for any j ? 0, any randomized learning algorithm L, any n - n(j), and any A - A(n; j), there is an example sequence S 2 S 0 (1; A; n) such that A ln n: We extend these results and obtain the following theorems. The corresponding upper bounds are given in Theorems 5 and 6. Theorem 9 For any deterministic learning algorithm L, any k - 1, any n - 2k, and any A - 0, there is an example sequence S 2 S 0 (k; A; n) such that Theorem 10 There are functions n(j) and A(n; j) such that for any j ? 0, any randomized learning algorithm L, any k - 1, any n - kn(j), and any A - kA(n; j), there is an example sequence S 2 S 0 (k; A; n) such that r Proof of Theorems 9 and 10. The proof is by a reduction to the case 1. The n attributes are divided into k groups G i , such that each group consists of attributes. Furthermore we choose numbers a i - \Xi A with A. For each group G i we choose a sequence S i accordingly to Lemmas 4 and 5, respectively, such that for any learning algorithm and a These sequences S i can be extended to sequences S 0 i with n attributes by setting all the attributes not in group i to 0. Concatenating the expanded sequences S 0 we get a sequence S. It is easy to see that S 2 S(k; A; n). On the other hand any learning algorithm for sequences with n attributes can be transformed into a learning algorithm for sequences with a smaller number of attributes by setting the missing attributes to 0. Thus on each subsequence S 0 of S learning algorithm L makes at least as many mistakes as given in (6) and (7), respectively. Hence and r s A r if the function A(n; j) is chosen appropriately. 2 The last theorem shows that for randomized algorithms the constant of 1 before A in Theorem 6 is optimal and that the best constant before the square root term is in [1; 2]. 9. Conclusion We developed algorithm Swin which is a variant of Winnow for on-line learning of disjunctions subject to target shift. We proved worst case mistake bounds for Swin which hold for any sequence of examples and any kind of target drift (where the amount of error in the examples and the amount of shifts is bounded). There is a deterministic and a randomized version of Swin where the analysis of the randomized version is more involved and interesting in its own right. Lower bounds show that our worst case mistake bounds are close to optimal in some cases. Computer experiments highlight that an explicit mechanism is necessary to make the algorithm adaptive to target shifts. Acknowledgments We would like to thank Mark Herbster and Nick Littlestone for valuable discussions. We also thank the anonymous referees for their helpful comments. Notes 1. By expanding the dimension to 2n, learning non-monotone disjunctions reduces to the monotone case. 2. Essentially one has to describe when a shift occurs and which literal is shifted. Obviously there is no necessity to shift if the current disjunction is correct on the current example. Thus only in some of the trials in which the current disjunction would make a mistake the disjunction is shifted. Since the target schedule might make up to A mistakes due to attribute errors and there are up to Z shifts, we get up to A + Z trials which are candidates for shifts. Choosing Z of them and choosing one literal for each shift gives Z possibilities. 3. For this potential function the weights must be positive. Negative weights are handled via a reduction (Littlestone, 1988, 1989; Haussler et al., 1994). 4. In the worst case the randomized algorithm makes a mistake with probability 1/2 in each trial and the deterministic algorithm always breaks the tie in the wrong way such that it makes a mistake in each trial. Thus the number of mistakes of the deterministic algorithm is twice the expected number of mistakes of the randomized algorithm. 5. Formally let r Such a sequence hr j i exists if p(\Delta) is not equal to 1 everywhere and if there is a value r with 1. For functions p(\Delta) not satisfying these conditions algorithm Swin can be forced to make an unbounded number of mistakes even in the absence of attribute errors. --R Tracking the best disjunction. How to use expert advice. Pattern classification and scene analysis. Tight worst-case loss bounds for predicting with expert advice (Tech Predicting nearly as well as the best pruning of a decision tree. Tracking the best expert. Entropy optimization principles with applications. Additive versus exponentiated gradient updates for linear prediction. The perceptron algorithm vs. linear vs. logarithmic mistake bounds when few input variables are relevant. Mistake bounds and logarithmic linear-threshold learning algorithms Redundant noisy attributes learning theory (pp. The weighted majority algorithm. Information and Computation Efficient learning with virtual threshold gates. learning theory (pp. --TR --CTR Mark Herbster , Manfred K. Warmuth, Tracking the Best Expert, Machine Learning, v.32 n.2, p.151-178, Aug. 1998 D. P. Helmbold , S. Panizza , M. K. Warmuth, Direct and indirect algorithms for on-line learning of disjunctions, Theoretical Computer Science, v.284 n.1, p.109-142, 6 July 2002 Chris Mesterharm, Tracking linear-threshold concepts with Winnow, The Journal of Machine Learning Research, 4, 12/1/2003 Mark Herbster , Manfred K. Warmuth, Tracking the best regressor, Proceedings of the eleventh annual conference on Computational learning theory, p.24-31, July 24-26, 1998, Madison, Wisconsin, United States Manfred K. Warmuth, Winnowing subspaces, Proceedings of the 24th international conference on Machine learning, p.999-1006, June 20-24, 2007, Corvalis, Oregon Gentile , Nick Littlestone, The robustness of the Olivier Bousquet , Manfred K. Warmuth, Tracking a small set of experts by mixing past posteriors, The Journal of Machine Learning Research, 3, 3/1/2003 Claudio Gentile, The Robustness of the p-Norm Algorithms, Machine Learning, v.53 n.3, p.265-299, December Giovanni Cavallanti , Nicol Cesa-Bianchi , Claudio Gentile, Tracking the best hyperplane with a simple budget Perceptron, Machine Learning, v.69 n.2-3, p.143-167, December 2007 Mark Herbster , Manfred K. Warmuth, Tracking the best linear predictor, The Journal of Machine Learning Research, 1, p.281-309, 9/1/2001 S. B. Kotsiantis , I. D. Zaharakis , P. E. Pintelas, Machine learning: a review of classification and combining techniques, Artificial Intelligence Review, v.26 n.3, p.159-190, November 2006 Peter Auer, Using confidence bounds for exploitation-exploration trade-offs, The Journal of Machine Learning Research, 3, 3/1/2003

concept drift;prediction;computational learning theory;amortized analysis;on-line learning

296382

Tracking the Best Expert.

We generalize the recent relative loss bounds for on-line algorithms where the additional loss of the algorithm on the whole sequence of examples over the loss of the best expert is bounded. The generalization allows the sequence to be partitioned into segments, and the goal is to bound the additional loss of the algorithm over the sum of the losses of the best experts for each segment. This is to model situations in which the examples change and different experts are best for certain segments of the sequence of examples. In the single segment case, the additional loss is proportional to log n, where n is the number of experts and the constant of proportionality depends on the loss function. Our algorithms do not produce the best partition; however the loss bound shows that our predictions are close to those of the best partition. When the number of segments is k+1 and the sequence is of length &ell;, we can bound the additional loss of our algorithm over the best partition by O(k \log n+k \log(&ell;/k)). For the case when the loss per trial is bounded by one, we obtain an algorithm whose additional loss over the loss of the best partition is independent of the length of the sequence. The additional loss becomes O(k\log n+ k \log(L/k)), where L is the loss of the best partitionwith k+1 segments. Our algorithms for tracking the predictions of the best expert aresimple adaptations of Vovk's original algorithm for the single best expert case. As in the original algorithms, we keep one weight per expert, and spend O(1) time per weight in each trial.

Introduction Consider the following on-line learning model. The learning occurs in a series of trials labeled In each trial t the goal is to predict the outcome y t 2 [0; 1] which is received at the end of the trial. At the beginning of trial t, the algorithm receives an n-tuple x t . The element x t;i 2 [0; 1] of the n-tuple x t represents the prediction of an expert E i of the value of the outcome y t on trial t. The algorithm then produces a prediction - based on the current expert prediction tuple x t , and on past predictions and outcomes. At the end of the trial, the algorithm receives the outcome y t . The algorithm then incurs a loss measuring the discrepancy between the prediction - y t and the outcome y t . Similarly, each expert incurs a loss as well. A possible goal is to minimize the total loss of the algorithm over all ' trials on an arbitrary sequence of instance outcome pairs (such pairs are called examples). Since we make no assumptions about the relationship between the prediction of experts and the outcome (y t ), there is always some sequence of y t that is * The authors were supported by NSF grant CCR-9700201. An extended abstract appeared in (Herbster & Warmuth, 1995) M. HERBSTER AND M. K. WARMUTH "far away" from the predictions - y t of any particular algorithm. Thus, minimizing the total loss over an arbitrary sequence of examples is an unreasonable goal. A refined relativized goal is to minimize the additional loss of the algorithm over the loss of the best expert on the whole sequence. If all experts have large loss then this goal might actually be easy to achieve, since for all algorithms the additional loss over the loss of the best expert may then be small. However, if at least one expert predicts well, then the algorithm must "learn" this quickly and produce predictions which are "close" to the predictions of the best expert in the sense that the additional loss of the algorithm over the loss of the best expert is bounded. This expert framework might be used in various settings. For example, the experts might predict the chance of rain or the likelihood that the stock market will rise or fall. Another setting is that the experts might themselves be various sub-algorithms for recognizing particular patterns. The "master" algorithm that combines the experts' predictions does not need to know the particular problem domain. It simply keeps one weight per expert, representing the belief in the expert's prediction, and then decreases the weight as a function of the loss of the expert. Previous work of Vovk (Vovk, 1998) and others (Littlestone & Warmuth, 1994; Haussler, Kivinen & Warmuth, 1998) has produced an algorithm for which there is an upper bound on the additional loss of the algorithm over the loss of the best expert. Algorithms that compare against the loss of the best expert are called Static-expert algorithms in this paper. The additional loss bounds for these algorithms have the form c ln n for a large class of loss functions, where c is a constant which only depends on the loss function L, and n is the number of experts. This class of loss functions contains essentially all common loss functions except for the absolute loss and the discrete loss 1 (counting prediction mistakes), which are treated as special cases (Littlestone & Warmuth, 1994; Vovk, 1995; Cesa-Bianchi, Freund, Haussler, Helmbold, Schapire & Warmuth, 1997). For example, if the loss function is the square or relative entropy loss, then respectively (see Section 2 for definitions of the loss functions). In the paper we consider a modification of the above goal introduced by Littlestone and Warmuth (Littlestone & Warmuth, 1994), in which the sequence of examples is subdivided into k segments of arbitrary length and distribution. Each segment has an associated expert. The sequence of segments and its associated sequence of experts is called a partition. The loss of a partition is the sum of the total losses of the experts associated with each segment. The best partition of size k is the partition with k segments that has the smallest loss. The modified goal is to perform well relative to the best partition of size k. This goal is to model real life situations where the "nature" of the examples might change and a different expert produces better predictions. For example, the patterns might change and different sub-algorithms may predict better for different segments of the on-line sequence of patterns. We seek to design master algorithms that "track" the performance of the best sequence of experts in the sense that they incur small additional loss over the best partition of size k. If the whole sequence of examples was given ahead of time, then one could compute the best partition of a certain size and the associated experts using dynamic programming. Our algorithms get the examples on-line and never produce the best partition. Even so, we are able to bound the additional loss over the best off-line partition for an arbitrary sequence of examples. When there are ' trials, k experts, there are distinct partitions. We can immediately get a good bound for this problem by expanding the set of n experts into experts." Each partition-expert represents a single partition of the trial sequence, and predicts on each trial as the expert associated with the segment which contains the current trial. Thus, using the Static-expert algorithm we obtain a bound of c ln of the additional loss of the algorithm over the loss of the best partition. There are two problems: first, the algorithm is inefficient, since the number of partition-experts is exponential in the number of partitions; and second, the bound on the additional loss grows with the sequence length. We were able to overcome both problems. Instead of keeping one weight for the exponentially many partitions, we can get away with keeping only one weight per expert, as done in the Static-expert algorithm. So the "tracking" of the predictions of the best partition is essentially for free. If there are n sub-algorithms or experts whose predictions we want to combine, then as in the Static-expert algorithm the new master algorithm takes only O(n) additional time per trial over the time required for simulating the n sub-algorithms. We develop two main algorithms: the Fixed-share Algorithm, and the Variable- share Algorithm. Both of these are based on the Static-expert algorithms which maintain a weight of the form e \GammajT i for each expert (cf. Littlestone & Warmuth, 1994; Vovk, 1995), where T i is total past loss of the expert i in past trials. In each trial the master algorithm combines the experts' predictions using the current weights of the experts. When the outcome of the trial is received, we multiply the weight of every expert i by e \GammajL i , where L i is the loss of expert i in the current trial. We call this update of the weights the Loss Update. We modify the Static-expert Algorithm by adding an additional update to obtain our algorithms. Since in our model the best expert may shift over a series of trials, we cannot simply use weights of the form e \GammajT i , because before an expert is optimal for a segment its loss in prior segments may be arbitrarily large, and thus its weight may become arbitrarily small. So we need to modify the Static-expert Algorithm so that small weights can be recovered quickly. For this reason, each expert "shares" a portion of its weight with the other experts after the Loss Update; we call this the Share Update. Both the Fixed-share and Variable-share Algorithm first do the Loss Update followed by a Share Update, which differs for each algorithm. In a Share Update, a fraction of each experts' weight is added to the weight of each other expert. In the Fixed-share Algorithm the experts share a fixed fraction of their weights with each other. This guarantees that the ratio of the weight of any expert to the total weight of all the experts may be bounded from below. Different forms of lower bounding the weights have been used by the Wml algorithm and in the companion paper for learning shifting disjunctions (Auer & Warmuth, 1998) that appears in this journal issue. The latter two methods have been applied to learning problems where the loss is the discrete 4 M. HERBSTER AND M. K. WARMUTH loss (i.e. counting mistakes). In contrast our methods work for the same general class of continuous loss functions that the Static-expert algorithms can handle (Vovk, 1998; Haussler et al., 1998). This class includes all common loss functions such as the square loss, the relative entropy loss, and the hellinger loss. For this class there are tight bounds on the additional loss (Haussler et al., 1998) of the algorithm over the loss of the best expert (i.e., the non-shifting case). The Fixed- share Algorithm obtains the additional loss of O(c[(k+1) log n+k log ' k +k]), which is essentially the same as the sketched algorithm that uses the Static-expert algorithm with exponentially many partition-experts. The salient feature of the Fixed-share Algorithm is that it still uses O(1) time per expert per trial. However, this algorithm's additional loss still depends on the length of the sequence. Our lower bounds give some partial evidence that this seems to be unavoidable for loss functions for which the loss in a single trial can be unbounded (such as for the relative entropy loss). For the case when the loss in a particular trial is at most one (such as for the square loss), we develop a second algorithm called the Variable- share Algorithm. This algorithm obtains bounds on the additional loss that are independent of the length of the sequence. It also shares weights after the Loss Update; however, the amount each expert shares now is commensurate with the loss of the expert in the current trial. In particular, when an expert has no loss, it does not share any weight. Both versions of our Share Update are trivial to implement and cost a constant amount of time for each of the n weights. Although the algorithms are easy to describe, proving the additional loss bounds takes some care. We believe that our techniques constitute a practical method for tracking the predictions of the best expert with provable worst-case additional loss bounds. The essential ingredient for our success in a non-stationary setting, seems to be an algorithm for the stationary setting with a multiplicative weight update whose loss bound grows logarithmically with the dimension of the problem. Besides Vovk's Aggregating Algorithm (Vovk, 1998) and the Weighted Majority Algorithm (Littlestone & Warmuth, 1994), which only use the Loss Update, and are the basis of this work, a number of such algorithms have been developed. Examples are algorithms for learning linear threshold functions (Littlestone, 1988; Littlestone, 1989), and algorithms whose additional loss bound over the loss of the best linear combination of experts or sigmoided linear combination of experts is bounded (Kivinen & Warmuth, 1997; Helmbold, Kivinen & Warmuth, 1995). Significant progress has recently been achieved for other non-stationary settings building on the techniques developed in this paper (see discussion in the Conclusion Section). The paper is outlined as follows. After some preliminaries (Section 2), we present the algorithms (Section 3), and give the basic proof techniques (Section 4). Sections 5 and 6 contain the detailed proofs for the Fixed-share and Variable- share algorithms, respectively. The absolute loss is treated as a special case in Section 7. Section 8 discusses a subtle but powerful generalization of the Variable- share Algorithm, called the Proximity-variable-share Algorithm. The generalization leads to improved bounds for the case when best expert of the next segment is always likely to be "close" to the previous expert. Some preliminary lower bounds are given in Section 9. Simulation results on artificial data that exemplify our methods are given in Section 10. Finally, in Section 11 we conclude with a discussion of recent work. The casual reader who might not be interested in the detailed proofs is recommended to read the sections containing the preliminaries (Section 2), the algorithms (Section 3) and the simulations (Section 10). 2. Preliminaries Let ' denote the number of trials and n denote the number of experts labeled When convenient we simply refer to an expert by its index; thus "expert i" refers to expert E i . The prediction of all n experts in trial t is referred to by the prediction tuple x t , while the prediction of expert i on trial t is denoted by x t;i : These experts may be viewed as oracles external to the algorithm, and thus may represent the predictions of a neural net, a decision tree, a physical sensor or perhaps even of a human expert. The outcome of a trial t is y t , while the prediction of the algorithm in trial t is - y t . The instance-outcome pair called the t-th example. In this paper the outcomes, the expert predictions and the predictions of the algorithm are all in [0; 1]. Throughout this paper S always denotes an arbitrary sequence of examples, i.e. any sequence of elements from [0; 1] n \Theta [0; 1] of any length '. A loss function L(p; q) is a function We consider four loss functions in this paper: the square, the relative entropy, the hellinger, and the absolute loss: ent (p; hel (p; On trial t the loss of the algorithm A is L(y Similarly, the loss of expert i on trial t is L(y t ; x t;i ). We call a subsequence of contiguous trials a segment. The notation non-negative integers t - t 0 denotes a segment starting on trial number t and ending on the trial t 0 . Rounded parens are used if the ending trial is not included in the segment. For the current sequence S we abbreviate the loss of expert i on the segment [t::t 0 ) by L([t::t 0 The loss of the algorithm A over the whole trial sequence S is defined as L(S; We are now ready to give the main definition of this paper that is used for scenarios in which the best expert changes over time. Informally a k-partition slices a sequence into k segments with an expert being associated with each segment. Formally, a k-partition, denoted by P ';n;k;t;e (S), consists of three positive integers '; n; k; and two tuples t and e of positive integers. The number ' is the length of the trial sequence S, n is the size of the expert pool, and k is number of target shifts '). The tuple t has k elements and Each t i refers to one of the ' trials, and by convention we use 1: The tuple t divides the trial sequence S into 6 M. HERBSTER AND M. K. WARMUTH Parameters: Initialization: Initialize the weights to w s t;i . Predict with Loss Update: After receiving the tth outcome y t , Share Updates of all three algorithms: Static-expert t;i "no Share Update" Fixed-share (4) Variable-share (5) Figure 1. The Static-expert, Fixed-share, and Variable-share algorithms called the ith segment. The 0th segment is also referred to as the initial segment. The tuple e has k+1 elements . The element e i denotes the expert which is associated with the ith segment [t i ::t i+1 ). The loss of a given k-partition for loss function L and trial sequence S is 3. The Algorithms There are four algorithms considered in this paper - Static-expert, Fixed- share, Variable-share and Proximity-variable-share. The first three are summarized in Figure 1. The Proximity-variable-share Algorithm is a generalization of the Variable-share Algorithm; this algorithm is given in Figure 3. The discussion of this generalization is deferred to Section 8. For all algorithms the learning process proceeds in trials, where t - 1 denotes the trial number. The algo- rithms maintain one positive weight per expert. The weight w s t;i (or its normalized version v s should be thought of as a measurement of the algorithm's belief in the quality of the ith expert's predictions at the start of trial t. The weight of each expert is initialized to 1=n. The algorithms have the following three parameters: j; c and ff. The parameter j is a learning rate quantifying how drastic the first update will be. The parameter c will be set to 1=j for most loss functions. (The absolute loss is an exception treated separately in Section 7.) The parameter ff quantifies the rate of shifting that is expected to occur. The Fixed-share Algorithm is designed for potentially unbounded loss functions, such as the relative entropy loss. The Variable-share Algorithm assumes that the loss per trial lies in [0; 1]. For the Fixed-share Al- gorithm, ff is the rate of shifting per trial. Thus, if five shifts are expected in a 1000 trial sequence, then 1=200. For the Variable-share Algorithm, ff is approximately the rate of shifting per unit of loss of the best partition. That is, if five shifts are expected to occur in a partition with a total loss of 80, then ff - 1=16. The tunings of the parameters j and c are considered in greater depth in Section 4, and for ff in sections 5 and 6. Finally, the Static-expert Algorithm does not use the parameter ff since it assumes that no shifting occurs. In each trial t the algorithm receives an instance summarizing the predictions of the n experts x t . The algorithm then plugs the current instance x t and normalized weights v t into the prediction function pred(v; x) in order to produce a prediction y t . In the simplest case, the algorithm predicts with the weighted mean of the experts' predictions, i.e., pred(v; more sophisticated prediction function introduced by Vovk (Vovk, 1998) will be discussed in Section 4. After predicting, the algorithm performs two update steps. The first update is the Loss Update; the second is the Share Update. In the Loss Update the weight of expert i is multiplied by e \GammajL i , where L i is the loss of the i-th expert in the current trial. Thus, no update occurs when L The learning rate j intensifies the effect of this update. We use w m t;i to denote the weights in the middle of the two updates. These weights will be referred to as intermediate weights. The Share Update for the Static-expert Algorithm is vacuous. However, for the other algorithms the Share Update is crucial. We briefly argue for the necessity of the share updates in the non-stationary setting, and then give an intuitive description of how they function. When we move from predicting as well as the best expert to predicting as well as a sequence of experts, the Loss Update is no longer appropriate as the sole update. Assume we have two experts and two segments. In the first segment Expert 1 has small loss and Expert 2 a large loss. The roles are reversed for the second segment. By the end of the first segment, the Loss Update has caused the weight of Expert 2 to be almost zero. However, during the second segment the predictions of Expert 2 are important, and its weight needs to be recovered quickly. The share updates make sure that this is possible. The simulation in Section 10 furthers the intuition for why the share updates are needed. The two share updates are summarized 8 M. HERBSTER AND M. K. WARMUTH below. A straightforward implementation costs O(n) time per expert per trial: Fixed-share: w s ff Variable-share: w s In contrast, the implementations in Figure 1, that use the intermediate variable "pool" cost O(1) time per expert per trial. After the Loss Update, every expert "shares" a fraction of its weight equally with every other expert. The received weight enables an expert to recover its weights quickly relative to the other experts. In the Fixed-share Update (6) each expert shares a fraction of ff of its weight in each trial. If one expert is perfect for a long segment, this type of sharing is not optimal, since the perfect expert keeps on sharing weight with possible non-perfect experts. The Variable-share Update (7) is more sophisticated: roughly, an expert shares weight when its loss is large. A perfect expert doesn't share, and if all other experts have high loss, it will eventually collect all the weight. However, when a perfect expert starts to incur high loss, it will rapidly begin to share its weight with the other experts, allowing a now good expert with previously small relative weight to recover quickly. As discussed above the parameter ff is the shifting rate. In the introduction we discussed an algorithm that uses exponentially many static experts, one for each partition. Our goal was to achieve bounds close to those of this inefficient algorithm by using only n weights. The bounds we obtain for our share algorithms are only slightly weaker than the partition-expert algorithm and gracefully degrade when neither the length of the sequence ' nor the number of shifts k are known in advance. 4. Prediction Functions and Proof Techniques We consider two choices of prediction functions. The simplest prediction is the weighted mean (Warmuth, 1997): pred wmean (v; A more sophisticated prediction function giving slightly better bounds was introduced by Vovk (Vovk, 1998; Haussler et al., 1998). Define L 0 z). Both of these functions must be monotone. Let L \Gamma1 1 (z) denote the inverses of L 0 (z) and L 1 (z). Vovk's prediction is now defined in two steps by pred Vovk (v; Loss c values: (j = 1=c) Functions: pred wmean (v; x) pred Vovk (v; x) ent (p; hel (p; q) 1 1= Figure 2. (c; 1=c)-realizability: c values for loss and prediction function pairings. The following definition is a technical condition on the relation between the prediction function pred(v; x), the loss function L, and the constants c and j. et al., 1998; Vovk, 1998) A loss function L and prediction function pred are (c; j)-realizable for the constants c and j if for all of total weight 1. We consider four loss functions in this paper: the square, the relative entropy, the hellinger, and the absolute loss (see Section 2). However, the algorithms are not limited to these loss functions. The techniques in (Vovk, 1998; Haussler et al., 1998; Warmuth, 1997) can determine the constants c and j for a wide class of loss functions. The algorithm is also easy to adapt for classification by using the majority vote (Littlestone & Warmuth, 1994) for the prediction function, and counting mistakes for the loss. In a practical application, no worst-case loss bounds may be provable for the given loss function. However, the share updates may still be useful. For an interesting application to the prediction of disk idle time see the work of Helmbold et al. (Helmbold, Long & Sherrod, 1996). The square, relative entropy and hellinger losses are (c; j)-realizable for both pred wmean and pred Vovk with (j = 1=c). The values of c (and hence of the two prediction functions are summarized in Figure 2. Since the absolute loss has more complex bounds, we treat it in a section of its own. A smaller value of c leads to a smaller loss bound (see Lemma 1). The c values for pred Vovk (cf. column two of Figure 2) are optimal for a large class of loss functions (Haussler et al., 1998). The proof of the loss bounds for each of the algorithms is based on the following lemma. The lemma embodies a key feature of the algorithms: the prediction is done such that the loss incurred by the algorithm is tempered by a corresponding change in total weight. This lemma gives the same inequality as the lemmas used in (Vovk, 1998; Haussler et al., 1998). The proof here is essentially the same, since the share updates do not change the total weight W t;i . Haussler et al., 1998) For any sequence of examples S and for any expert i, the total loss of the master algorithms in Figure 1 may be M. HERBSTER AND M. K. WARMUTH bounded by when the loss function L and prediction function pred is (c; j)-realizable (cf. Definition 1 and Figure 2). Proof: Since L and pred are (c; j)-realizable, we have by Definition 1 that Since the share updates do not change, the total weight t;i is W t+1 . This implies that Hence, since W So far we have used the same basic technique as in (Littlestone & Warmuth, 1994; Vovk, 1995; Cesa-Bianchi et al., 1997; Haussler et al., 1998), i.e., c ln W t becomes the potential function in an amortized analysis. In the static expert case (when 1=c) the final weights have the form w s =n. Thus the above lemma leads to the bound relating the loss of the algorithm to the loss of any static expert. The share updates make it much more difficult to lower bound the final weights. Intuitively, there has to be sufficient sharing so that the weights can recover quickly. However, there should not be too much sharing, so that the final weights are not too low. In the following sections we bound final weights of individual experts in terms of the loss of a partition. The loss of any partition (L(P ';n;k;t;e (S))) is just the sum of the sequence of losses defined by the sequence of experts in the partition. When an expert accumulates loss over a segment, we bound its weight using Lemma 2 for the Fixed-share Algorithm and Lemma 7 for the Variable- share Algorithm. Since a partition is composed of distinct segments, we must also quantify how the weight is transferred from the expert associated with a segment to the expert associated with the following segment; this is done with Lemma 3 for the Fixed-share Algorithm and Lemma 8 for the Variable-share Algorithm. The lower bounds on the weights are then combined with Lemma 1 to bound the total loss of the Fixed-share Algorithm (Theorem 1) and the Variable-share Algorithm (Theorem 2). 5. Fixed-share Analysis This algorithm works for unbounded loss functions, but its total additional loss grows with the length of the sequence. Lemma 2 For any sequence of examples S the intermediate weight of expert i on trial t 0 is at least e \GammajL([t::t 0 ];i) times the weight of expert i at the start of trial t, where t - t 0 . Formally we have Proof: The combined Loss and Fixed-share Update (Equation (6)) can be rewritten as Then if we drop the additive term produced by the Share Update, we have We apply the above iteratively on the trials [t::t 0 ). Since we are bounding w m weights in trial t 0 after the Loss Update), the weight on trial t 0 is only reduced by a factor of e \GammajL(y t 0 ;x t 0 ;i ) . Therefore we have Y r=t By simple algebra and the definition of L([a::b]; i) the bound of the lemma follows. Lemma 3 For any sequence of examples S, the weight of an expert i at the start of trial 1 is at least ff times the intermediate weight of any other expert j on trial t. Proof: Expanding the Fixed-share Update (4) we have Thus w s ff and we are done. We can now bound the additional loss. M. HERBSTER AND M. K. WARMUTH Theorem 1 Let S be any sequence of examples and let L and pred be (c; j)- realizable. Then for any k-shift sequence partition P ';n;k;t;e (S) the total loss of the Fixed-share Algorithm with parameter ff satisfies Proof: Recall that e k is the expert of the last segment. By Lemma 1, with we have We bound w s '+1;ek by noting that it "follows" the weight in an arbitrary partition. This is expressed in the following telescoping product: Y Thus, applying lemmas 3 and 2, we have Y The final term w s equals one, since we do not apply the Share Update on the final trial; therefore by the definition of L(P ';n;k;t;e (S)), we have e \GammajL(P ';n;k;t;e We then substitute the above bound on w s '+1;ek into (16) and simplify to obtain (15). The bound of Theorem 1 holds for all k, and there is a tradeoff between the terms ck ln n and cjL(P ';n;k;t;e (S)); i.e., when k is small the ck ln n term is small and the cjL(P ';n;k;t;e (S)) term is large, and vice-versa. The optimal choice of ff (obtained by differentiating the bound of Theorem 1) is ff . The following corollary rewrites the bound of Theorem 1 in terms of the optimal parameter choice ff . The corollary gives an interpretation of the theorem's bound in terms of code length. We introduce the following notation. Let 1\Gammap be the binary entropy measured in nats, and 1\Gammaq be the binary relative entropy in nats. 2 Corollary 1 Let S be any sequence of examples and let L and pred be (c; j)- realizable. Then for any k-shift sequence partition P ';n;k;t;e (S) the total loss of the Fixed-share Algorithm with parameter ff satisfies ck where ff . When , then this bound becomes For the interpretation of the bound we ignore the constants c, j and the difference between nats and bits. The terms ln n and k ln(n \Gamma 1) account for encoding the experts of the partition: log n bits for the initial expert and log(n \Gamma 1) bits for each expert thereafter. Finally, we need to encode where the k shifts occur (the inner boundaries of the partition). If ff is interpreted as the probability that a shift occurs on any of the '\Gamma1 trials, then the term ('\Gamma1) [H(ff ) +D(ff kff)] corresponds to the expected optimal code length (see Chapter 5 of (Cover & Thomas, 1991)) if we code the shifts with the estimate ff instead of the true probability ff . This bound is thus an example of the close similarity between prediction and coding as brought out by many papers (e.g. (Feder, Merhav & Gutman, 1992)). Note that the ff that minimizes the bound of Theorem 1 depends on k and ' which are unknown to the learner. In practice a good choice of ff may be determined experimentally. However, if we have an upper bound on ' and a lower bound on k we may tune ff in terms of these bounds. Corollary 2 Let S be any sequence of examples and - ' and - k be any positive integers such that - 1. Then by setting 1), the loss of the Fixed- share Algorithm can be bounded by where P ';n;k;t;e (S) is any partition of S such that ' - ' and k - k. Proof: Recall the loss bound given in Theorem 1. By setting , we have We now separate out the term apply the inequality The last inequality follows from the condition that ' - ' and - We obtain the bound of the corollary by replacing in Equation (20) by its upper bound - k. 3 14 M. HERBSTER AND M. K. WARMUTH 6. Variable-share analysis The Variable-share algorithm assumes that the loss of each expert per trial lies in [0; 1]. Hence the Variable-share Algorithm works in combination with the square, hellinger, or absolute loss functions but not with the relative entropy loss function. The Variable-share Algorithm has an upper bound on the additional loss of the algorithm which is independent of the length of the trial sequence. We will abbreviate w s t;i with w t;i , since in this section we will not need to refer to the weight of an expert in the middle of a trial. We first give two technical lemmas that follow from convexity in r of fi r . c). Applying the first inequality of Lemma 4 to the RHS we have c + db d - b 1\Gammac , and thus Lemma 6 At the beginning of trial t +1, we may lower bound the weight of expert i by either Expression (a) or Expression (b), where j is any expert different from i: ae w t;i e \GammajL(y t ;x t;i Proof: Expanding the Loss Update and the Variable-share Update for a trial (cf. (7)) we have Expression (a) is obtained by dropping the summation term. For Expression (b) we drop all but one summand of the second term: w . We then apply Lemma 4 and obtain (b). Lemma 7 The weight of expert i from the start of trial t to the start of trial t 0 , reduced by no more than a factor of [e \Gammaj \Theta Proof: From Lemma 6(a), we have that on trial t the weight of expert i is reduced as follows: w t+1;i we apply this iteratively on the Y r=t \Theta In Lemma 6(b) we lower bound the weight transferred from expert p to expert q in a single trial. In the next lemma we show how weight is transferred over a sequence of trials. Lemma 8 For any distinct experts p and q, if L([t::t 0 2, then on trial t may lower bound the weight of expert q by - \Theta e \Gammaj Proof: As expert p accumulates loss in trials t::t 0 , it transfers part of its weight to the other specifically to expert q, via the Variable-share Update. Let a i , for t - i - t 0 , denote the weight transferred by expert p to expert q in trial i=t a i denote the total weight transferred from expert p to expert q in trials [t::t 0 ]: The transferred weight, however, is still reduced as a function of the loss of expert q in successive trials. By Lemma 7, the weight a i added in trial i is reduced by a factor of [e \Gammaj during a i \Theta We lower bound each factor [e \Gammaj by [e \Gammaj , and thus \Theta e \Gammaj To complete the proof of the lemma we still need to lower bound the total transferred weight A by w t;p ff l i be the loss of expert p on trial i, i.e. From our assumption, we have 1 - 2. By direct application of Lemma 6(b), the weight a t transferred by expert p to expert q in the first trial t of the segment is at least w t;p ff l t e \Gammajl t . Likewise, we apply Lemma 7 over trials [t::i) to expert p, and then apply Lemma 6(b) on trial i. This gives us a lower bound for the transferred weights a i and the total transferred weight A: a ff a ff l M. HERBSTER AND M. K. WARMUTH We split the last sum into two terms: ff l ff We upper bound all exponents of (1 \Gamma ff) by one; we also replace the sum in the first exponent by its upper bound, . The substitutions 1, and then lead to an application of Lemma 5. Thus we rewrite the above inequality as ff \Theta ff and then apply Lemma 5. This gives us ff The proof of the loss bound for the Variable-share Algorithm proceeds analogously to the proof of the Fixed-share Algorithm's loss bound. In both cases we "follow" the weight of a sequence of experts along the sequence of segments. Within a segment we bound the weight reduction of an expert with Lemma 2 for the Fixed-share analysis and Lemma 7 for Variable-share analysis. When we pass from one segment to the next, we bound the weight of the expert corresponding to the new segment by the weight of the expert in the former segment with lemmas 3 and 8, respectively. The former lemma used for the Fixed- share Algorithm is very simple, since in each trial each expert always shared a fixed fraction of its weight. However, since the weight was shared on every trial, this produced a bound dependent on sequence length. In the Variable-share Algorithm we produce a bound independent of the length. This is accomplished by each expert sharing weight in accordance to its loss. However, if an expert does not accumulate significant loss, then we cannot use Lemma 8 to bound the weight of the following expert in terms of the previous expert. Nevertheless, if the former expert does not make significant loss in the current segment, this implies that we may bound the current segment with the former expert by collapsing the segments together. In other words, the collapsing of two consecutive segments ([t creates a single segment ([t which is associated with the expert of the first segment of the original two consecutive segments. We can do this for any segment; thus we determine our bound in terms of the related collapsed partition whose loss is not much worse. Lemma 9 For any partition P ';n;k;t;e (S) there exists a collapsed partition P ';n;k 0 such that for each segment (except the initial segment), the expert associated with the prior segment incurs at least one unit of loss, and the loss on the whole sequence of the collapsed partition exceeds the loss of the original partition by no more than the following properties hold: Proof: Recall that e i is the expert associated with the ith segment, which is comprised of the trials [t i ::t i+1 If in any segment i, the loss of the expert e associated with the prior segment (i \Gamma 1) is less than one, then we merge segment segment i. This combined segment in the new partition is associated with expert e i\Gamma1 . Formally in each iteration, we decrement k by one, and we delete e i and t i from the tuples e and t. We continue until (24) holds. We bound the loss of the collapsed partition P ';n;k 0 by noting that the loss of the new expert on the subsumed segment is at most one. Thus per application of the transformation, the loss increases by at most one. Thus since there are applications, we are done. Theorem 2 4 Let S be any sequence of examples, let L and pred be (c; j)-realizable, and let L have a [0,1] range. Then for any partition P ';n;k;t;e (S) the total loss of the Variable-share algorithm with parameter ff satisfies Proof: By Lemma 1 with Let P ';n;k;t;e (S) be an arbitrary partition. For this proof we need the property that the loss in each segment (except the initial segment), with regard to the expert associated with the prior segment, is at least one (cf (24)). If this property does not hold, we use Lemma 9 to replace P ';n;k;t;e (S) by a collapsed partition P ';n;k 0 for which the property does hold. If the property holds already for P ';n;k;t;e (S), then for notational convenience we will refer to P ';n;k;t;e (S) by P ';n;k 0 Recall that the loss of exceeds the loss of P ';n;k;t;e (S) by no more than Since (24) holds, there exists a trial q i in the ith segment (for 1 that L([t 0 1. We now express w '+1;e 0 as the telescoping product Y Applying lemmas 7 and 8 we have ii ff M. HERBSTER AND M. K. WARMUTH which simplifies to the following bound: ff The last inequality follows from (25). Thus if we substitute the above bound on simplify, we obtain the bound of the theorem. Again we cannot optimize the above upper bound as a function of ff, since k and L(P ';n;k;t;e (S)) are not known to the learning algorithm. Below we tune ff based on an upper bound of L(P ';n;k;t;e (S)). The same approach was used in Corollary 2. 5 Corollary 3 Let S be any sequence of examples and - L and - k be any positive reals. Then by setting L , the loss of the Variable-share Algorithm can be bounded as follows: ck where P ';n;k;t;e (S) is any partition such that L(P ';n;k;t;e (S)) - L, and in addition L. For any partition P ';n;k;t;e (S) for which L(P ';n;k;t;e (S)) - L, we obtain the upper bound ck Proof: We proceed by upper bounding the three terms containing ff from the bound of Theorem 2 (we use We rewrite the above as: We apply the identity ln(1 x and bound L(P ';n;k;t;e (S)) by - L, giving the following upper bound of the previous expression: L, then - L. Therefore the above is upper bounded by Using this expression to upper bound Equation (29), we obtain Equation (27). When L, we upper bound Equation (29) by The first term is bounded by 1- k. The second term (2 - k+ - is at most k ln 9 2 in the region thus the above is upper bounded by We use the above expression to upper bound Equation (29). This gives us Equation (28) and we are done. 7. Absolute Loss Analysis The absolute loss function L abs (p; j)-realizable with both the prediction functions pred Vovk and pred wmean ; however, cj ? 1. Thus the tuning is more complex, and for the sake of simplicity we use the weighted mean prediction (Littlestone & Warmuth, 1994) in this section. Theorem 3 (Littlestone & Warmuth, 1994) For 1), the absolute loss function L abs (p; j)-realizable for the prediction function pred wmean (v; x). To obtain a slightly tighter bound we could also have used the Vee Algorithm for the absolute loss, which is ((2 j)-realizable (Haussler et al., 1998). This algorithm takes O(n log n) time to produce its prediction. Both the weighted mean and the Vee prediction allow the outcomes to lie in [0; 1]. For binary outcomes with the absolute loss, O(n) time prediction functions exist with the same realizability criterion as the Vee prediction (Vovk, 1998; Cesa-Bianchi et al., 1997). Unlike the (c; 1=c)-realizable loss functions discussed earlier (cf. Figure 2), the absolute value loss does not have constant parameters, and thus it must be tuned. In practice, the tuning of j may be produced by numerical minimization of the upper bounds. However, we use a tuning of j produced by Freund and Schapire Theorem 4 (Lemma 4 P and Q. Q), where M. HERBSTER AND M. K. WARMUTH We now use the above tuning in the bound for the Variable-share Algorithm (Theorem 2). Theorem 5 Let the loss function be the absolute loss. Let S be any sequence of examples, and - L and - k be any positive reals such that k - k, L(P ';n;k;t;e (S)) - L, and k - L. Set the two parameters of the Variable-share algorithm ff and j to - k respectively, where - k and - k. Then the loss of the Algorithm with weighted mean prediction can be bounded as follows: Alternatively, let - L and - k be any positive reals such that k - k, L(P ';n;k;t;e (S)) - L, and k - L. Set the two parameters of the Variable-share algorithm ff and j to - k respectively, where - k and - k. Then the loss of the Algorithm with weighted mean prediction can be bounded as follows: 8. Proximity-variable-share Analysis In this section we discuss the Proximity-variable-share Algorithm (see Figure 3). Recall that in the Variable-share Algorithm each expert shared a fraction of weight dependent on its loss in each trial; that fraction is then shared uniformly among the remaining experts. The Proximity-variable-share Algorithm enables each expert to share non-uniformly to the other experts. The Proximity-variable-share Update now costs O(n) per expert per trial instead of O(1) (see Figure 3). This algorithm allows us to model situations where we have some prior knowledge about likely pairs of consecutive experts. Let us consider the parameters of the algorithm. The n-tuple - contains the initial weights of the algorithm, i.e., w s . The Parameters: Initialization: Initialize the weights to w s n . t;i . Predict with Loss Update: After receiving the tth outcome y t , Proximity-variable-share Update Figure 3. The Proximity-variable-share algorithm second additional parameter besides j and c is a complete directed graph - of size n without loops. The edge weight - j;k is the fraction of the weight shared by expert j to expert k. Naturally, for any vertex, all outgoing edges must be nonnegative and sum to one. The - 0 probability distribution is a prior for the initial expert and the probability distribution is a prior for which expert will follow expert j. Below is the upper bound for the Proximity-variable-share Algorithm. The Fixed- share Algorithm could be generalized similarly to take proximity into account. Theorem 6 Let S be any sequence of examples, let L and pred be (c; j)-realizable, and let L have a [0,1] range. Then for any partition P ';n;k;t;e (S), the total loss of the Proximity-variable-share Algorithm with parameter ff satisfies Proof: We omit the proof of this bound since it is similar to the corresponding proof of Theorem 2 for the Variable-share Algorithm: The only change is that the 1 fractions are replaced by the corresponding - parameters. Note that setting - gives the previous bound for the Variable-share Algorithm (Theorem 2). In that case the last sum is O(k ln n), accounting for the code length of the names of the best experts (except the first one). Using the Proximity-variable-share Algorithm we can get this last sum to O(k) in some cases. 22 M. HERBSTER AND M. K. WARMUTH For a simple example, assume that the processors are on a circular list and that for the two processors of distance d from processor i, - i;i+d mod 1=d 2 . Now if the next best expert is always at most a constant away from the previous one, then the last sum becomes O(k). Of course, other notions of closeness and choices of the - parameters might be suitable. Note that there is a price for decreasing the last sum: the update time is now O(n 2 ) per trial. However, if for each expert i all arrows that end at i are labeled with the same value, then the Share Update of the Proximity-variable-share Algorithm is still O(n). 9. Lower Bounds The upper bounds for the Fixed-share Algorithm grow with the length of the sequence. The additional loss of the algorithm over the loss of the best k-partition is approximately This holds for unbounded loss functions such as the relative entropy loss. When restricting the loss to lie in [0; 1], the Variable-share Algorithm gives an additional loss bound of approximately is the loss of the best k-partition and k ! L. One natural question is whether a similar reduction is possible for unbounded loss functions. In other words, whether for an unbounded loss function a bound of the same form is possible with ' replaced by minf'; Lg. We give evidence to the contrary. We give an adversary argument that forces any algorithm to make loss over the best one-partition (for which the adversary sets In this section we limit ourselves to giving this construction. It can easily be extended to an adversary that forces ln(n)+ln(' \Gamma log 2 n) additional loss over the best one-partition with n experts. By iterating the adversary, we may force additional loss over the best k-partition. (Here we assume that log 2 (n \Gamma 1) and ' are positive integers, and ' Theorem 7 For the relative entropy loss there exists an example sequence S of length ' with two experts such that L(P ';2;1;t;e there is a partition with a single shift of loss 0, and furthermore, for any algorithm A, Proof: The adversary's strategy is described in Figure 4. We use - y t to denote the prediction of an arbitrary learning algorithm, and L the loss at trial t. For convenience we number the trials from of There are two experts; one always predicts 0 and the other always predicts 1. The adversary returns a sequence of 0 outcomes followed by a sequence of 1 outcomes such that neither sequence is empty. Thus, there is a single shift in the best partition, and this partition has loss 0. 2. On trial 2 and to 1 otherwise. (Assume without loss of generality that - 2 and thus y 3. New trial: 4. If - '\Gammat then 5. else Go to step 7. then go to step 3. 7. Let y for the remaining trial(s) and exit. Figure 4. Adversary's strategy We now prove that L(S; thus proving the lemma. Clearly loss of generality assume - y . Note that the threshold for - y t is 1 '\Gammat . Furthermore, L ent (0; 1 and L ent (1; 1 t). Thus the conditions 4(i) and 5(i) follow. Condition 4(ii) holds by simple induction. If a shift occurs, then Condition 5(ii) holds, since by Condition 4(ii) in we have that '\Gammat . Therefore, when we add L t , which is at least ln(' \Gamma t) by Condition 5(i), we obtain Condition 5(ii) and we are done. If Step 5 is never executed then the shift to y occurs in the last trial Step 6 is skipped. Thus if Step 5 is never executed then in trial (Condition 4(ii)), which is again the bound of the lemma. We first reason that this lower bound is tight by showing that the upper bounds of the algorithms discussed in this paper are close to the lower bound. The number of partitions when 1). Thus we may expand the set of experts into partition-experts as discussed in the introduction. Using the Static-expert Algorithm with the weighted mean prediction gives an upper bound of on the total loss of the algorithm when the loss of the best partition is zero. This matches the above lower bound. Second, the bound of the Fixed-share Algorithm (cf. Corollary 1) is larger than the lower bound by '\Gamma2 ), and this additional term may be upper bounded by 1. M. HERBSTER AND M. K. WARMUTH total loss of the algorithms trials Loss of Variable Share Algorithm Loss of Static Algorithm (Vovk) Loss of Fix Share Algorithm Loss of typical expert Loss of best partition (k=3) Variable Share Loss Bound Fix Share Loss Bound Figure 5. Loss of the Variable-share Algorithm vs the Static-expert Algorithm scaled weights Figure 6. Relative Weights of the Variable-share Algorithm 10. Simulation Results In this section we discuss some simulations on artificial data. These simulations are mainly meant to provide a visualization of how our algorithms track the predictions of the best expert and should not be seen as empirical evidence of the practical usefulness of the algorithms. We believe that the merits of our algorithms are more clearly reflected in the strong upper bounds we prove in the theorems of 8000.10.30.50.70.9scaled weights trials Vovk Relative Weights Figure 7. Relative Weights of the Static-expert Algorithm the earlier sections. Simulations only show the loss of an algorithm for a typical sequence of examples. The bounds of this paper are worst-case bounds that hold even for adversarially-generated sequences of examples. Surprisingly, the losses of the algorithms in the simulations with random sequences are very close to the corresponding worst-case bounds which we have proven in this paper. Thus our simulations show that our loss bounds are tight for some sequences. We compared the performance of the Static-expert Algorithm to the two Share algorithms in the following setting. We chose to use the square loss as our loss function, because of its widespread use and because the task of tuning the learning rate for this loss function is simple. We used the Vovk prediction function (cf. Equation 9), and we chose accordance with Figure 2. We considered a sequence of 800 trials with four distinct segments, beginning at trials 1, 201, 401, and 601. On each trial the outcome (y t ) was 0. The prediction tuple contained the predictions of 64 experts. When we generated the predictions of the 64 experts, we chose a different expert as the best one for each segment. The best experts always have an expected loss of 1=120 per trial. The other 63 experts have an expected loss of 1=12 per trial. At the end of each segment a new "best expert" was chosen. Since the outcome was always 0, we generated these expected losses by sampling predictions from a uniform random distribution on (0; 1 for the "typical" and "best" experts, respectively. Thus the expected loss for the best 6 partition, denoted by the segment boundaries above, is 800 with a variance of oe 2 - :044. The actual loss of the best partition in the particular simulation used for the plots was 6:47. For the Fixed-share Algorithm we tuned f based on the values of using the ff f 26 M. HERBSTER AND M. K. WARMUTH tuning suggested in Corollary 1. For the Variable-share Algorithm we tuned ff v based on the values of using the ff v tuning suggested in Corollary 3. Using theorems 1 and 2 we calculated a worst case upper bound on the loss of the Fixed-share Algorithm and the Variable-share Algorithm of 24:89 and 21:50, respectively (see "\Theta" and "+" marks in Figure 5). The simulations on artificial data show that our worst-case bounds are rather tight even on this very simple artificial data. There are many heuristics for finding a suitable tuning. We used the tunings prescribed by our theorem, but noticed that for these types of simulations the results are relatively insensitive to the tuning of ff. For example, in calculating ff v for the Variable-share Algorithm when - was overestimated by 10 standard deviations, the loss bound for our algorithm increased by only 0:02, while the actual loss of the algorithm in the simulation increased by 0:17. In Figure 5, we have plotted the loss of the Static-expert Algorithm versus the loss of the two Share algorithms. Examination of the figure shows that on the first segment the Static-expert Algorithm performed comparably to the Share algo- rithms. However, on the remaining three segments, the Static-expert Algorithm performed poorly, in that its loss is essentially as bad as the loss of a "typical" expert (the slope of the total loss of a typical expert and the Static-expert Algorithm is essentially the same for the later segments). The Share algorithms performed poorly at the beginning of a new segment; however, they quickly "learned" the new "best" expert for the current segment. The Share algorithms' loss plateaued to almost the same slope as the slope of the total loss of the best expert. The two Share algorithms had the same qualitative behavior, even though the Fixed-share Algorithm incurred approximately 10% additional loss over the Variable-share Algorithm. In our simulations we tried learning rates j slightly smaller than two, and verified that even with other choices for the learning rates, the total loss of the Static-expert algorithm does not improve significantly. In Figures 6 and 7, we plotted the weights of the normalized weight vector w t that is maintained by the Variable-share Algorithm and the Static-expert Algorithm over the trial sequence. In Figure 6, we see that the Variable-share Algorithm shifts the relative weights rapidly. During the latter part of each segment, the relative weight of the best expert is almost one (the corresponding plot of the Fixed-share Algorithm is similar). On the other hand, we see in Figure 7 that the Static-expert Algorithm also "learned" the best expert for segment 1. However, the Static-expert Algorithm is unable to shift the relative weight sufficiently quickly, i.e. it takes the length of the second segment to partially "unlearn" the best expert of the first segment. The relative weights of the best experts for segments one and two essentially perform a "random walk" during the third segment. In the final segment, the relative weight of the best expert for segment three also performs a "random walk." In summary, we see these simulations as evidence that the Fixed-share and Variable-share Updates are necessary to track shifting experts. 11. Conclusion In this paper, we essentially gave a reduction for any multiplicative update algorithm that works well compared to the best expert for arbitrary segments of examples, to an algorithm that works well compared to the best partition, i.e. a concatenation of segments. Two types of share updates were analyzed. The Fixed- share Algorithm works well when the loss function can be unbounded, and the Variable-share Algorithm is suitable for the case when the range of the loss lies in [0,1]. The first method is essentially the same as the one used in the Wml algorithm of (Littlestone & Warmuth, 1994) and a recent alternate developed in (Auer Warmuth, 1998) for learning shifting disjunctions. When the loss is the discrete loss (as in classification problems), then these methods are simple and effective if the algorithm only updates after a mistake occurs (i.e., conservative updates). Our second method, the Variable-share Update, is more sophisticated. In particular, if one expert predicts perfectly for a while, then it can collect all the weight. However, if this expert is starting to incur large loss, then it shares weight with the other experts, helping the next best expert to recover its weight from zero. The methods presented here and in (Littlestone &Warmuth, 1994) have inspired a number of recent papers. Auer and Warmuth (1998) adapted the Winnow algorithm to learn shifting disjunctions. Comparing against the best shifting disjunction is more complicated than comparing against the best expert. However, since this is a classification problem a simple Sharing Update similar to the Fixed-share Update is sufficient. Our focus in this paper was to track the prediction of the best expert for the same class of loss functions for which the original Static-expert Algorithm of Vovk was developed (Vovk, 1998; Haussler et al., 1998). Our share updates have been applied experimentally for predicting disk idle times (Helmbold et al., 1996) and for the on-line management of investment portfolios (Singer, 1997). In addition, a reduction has been shown between expert and metrical task systems algorithms (Blum & Burch, 1997). The Share Update has been used successfully in the new domain of metrical task systems. A natural probabilistic interpretation of the Share algorithms has recently been given in (Vovk, 1997). In any particular application of the Share algorithms, it is necessary to consider how to choose the parameter ff. Theoretical techniques exist for the Fixed-share Algorithm for eliminating the need to choose the value of ff ahead of time. One method for tuning parameters (among other things) is the "specialist" framework of (Freund, Schapire, Singer & Warmuth, 1997), even though the bounds produced this way are not always optimal. Another method incorporates a prior distribution on all possible values of ff. For the sake of simplicity we have not discussed these methods (Herbster, 1997; Vovk, 1997; Singer, 1997) in this paper. 28 M. HERBSTER AND M. K. WARMUTH Acknowledgments We would like to thank Peter Auer, Phillip Long, Robert Schapire, and Volodya Vovk for valuable discussions. We also thank the anonymous referees for their helpful comments. Notes 1. The discrete loss is defined to be ae 2. Note that Lent (p; q). We use the D(pkq) notation here as is customary in information theory. 3. If we replace the assumption that k - k by 2 - k - ', we obtain a bound where the final term c - k is replaced by 2c - k: 4. Vovk has recently proved a sharper bound for this algorithm (Vovk, 1997): ff 5. Unlike Corollary 2 we do not need a lower bound on k. 6. We call the partition described by the segment boundaries 1, 201, 401, and 601, the best partition with respect to the tradeoff between k and L(P ';n;k;t;e (S)), as expressed implicitly in Theorem 2. --R Tracking the best disjunction. How to use expert advice. Elements of Information Theory. Universal prediction of individual sequences. IEEE Transactions on Information Theory A decision-theoretic generalization of on-line learning and an application to boosting Using and combining predictors that specialize. Sequential prediction of individual sequences under general loss functions. A dynamic disk spin-down technique for mobile computing Tracking the best expert II. Additive versus exponentiated gradient updates for linear prediction. Learning when irrelevant attributes abound: A new linear-threshold algorithm Mistake Bounds and Logarithmic Linear-threshold Learning Algorithms PhD thesis The weighted majority algorithm. Towards realistic and competitive portfolio selection algorithms. A game of prediction with expert advice. Derandomizing stochastic prediction strategies. Predicting with the dot-product in the experts framework --TR --CTR Atsuyoshi Nakamura, Learning specialist decision lists, Proceedings of the twelfth annual conference on Computational learning theory, p.215-225, July 07-09, 1999, Santa Cruz, California, United States Jeremy Z. Kolter , Marcus A. Maloof, Using additive expert ensembles to cope with concept drift, Proceedings of the 22nd international conference on Machine learning, p.449-456, August 07-11, 2005, Bonn, Germany V. Vovk, Probability theory for the Brier game, Theoretical Computer Science, v.261 n.1, p.57-79, 06/17/2001 Peter Auer , Manfred K. Warmuth, Tracking the Best Disjunction, Machine Learning, v.32 n.2, p.127-150, Aug. 1998 Olivier Bousquet , Manfred K. Warmuth, Tracking a small set of experts by mixing past posteriors, The Journal of Machine Learning Research, 3, 3/1/2003 Avrim Blum , Carl Burch, On-line Learning and the Metrical Task System Problem, Machine Learning, v.39 n.1, p.35-58, April 2000 Chris Mesterharm, Tracking linear-threshold concepts with Winnow, The Journal of Machine Learning Research, 4, 12/1/2003 Peter Auer, Using confidence bounds for exploitation-exploration trade-offs, The Journal of Machine Learning Research, 3, 3/1/2003 Giovanni Cavallanti , Nicol Cesa-Bianchi , Claudio Gentile, Tracking the best hyperplane with a simple budget Perceptron, Machine Learning, v.69 n.2-3, p.143-167, December 2007 Marco Barreno , Blaine Nelson , Russell Sears , Anthony D. Joseph , J. D. Tygar, Can machine learning be secure?, Proceedings of the 2006 ACM Symposium on Information, computer and communications security, March 21-24, 2006, Taipei, Taiwan Wei Yan , Christopher D. Clack, Diverse committees vote for dependable profits, Proceedings of the 9th annual conference on Genetic and evolutionary computation, July 07-11, 2007, London, England Mark Herbster , Manfred K. Warmuth, Tracking the best regressor, Proceedings of the eleventh annual conference on Computational learning theory, p.24-31, July 24-26, 1998, Madison, Wisconsin, United States Claudio Gentile, The Robustness of the p-Norm Algorithms, Machine Learning, v.53 n.3, p.265-299, December Mark Herbster , Manfred K. Warmuth, Tracking the best linear predictor, The Journal of Machine Learning Research, 1, p.281-309, 9/1/2001 Amol Deshpande , Zachary Ives , Vijayshankar Raman, Adaptive query processing, Foundations and Trends in Databases, v.1 n.1, p.1-140, January 2007

experts;multiplicative updates;shifting;amortized analysis;on-line learning

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On Asymptotics in Case of Linear Index-2 Differential-Algebraic Equations.

Asymptotic properties of solutions of general linear differential-algebraic equations (DAEs) and those of their numerical counterparts are discussed. New results on the asymptotic stability in the sense of Lyapunov as well as on contractive index-2 DAEs are given. The behavior of the backward differentiation formula (BDF), implicit Runge--Kutta (IRK), and projected implicit Runge--Kutta (PIRK) methods applied to such systems is investigated. In particular, we clarify the significance of certain subspaces closely related to the geometry of the DAE. Asymptotic properties like A-stability and L-stability are shown to be preserved if these subspaces are constant. Moreover, algebraically stable IRK(DAE) are B-stable under this condition. The general results are specialized to the case of index-2 Hessenberg systems.

Introduction . The present paper is devoted to the study of asymptotic properties of solutions of differential-algebraic equations (DAE's) on infinite intervals and those of their numerical counterparts in integration methods. It is rather surprising that, in spite of numerous papers on numerical integration, there are very few results in this respect. For index-1 DAE's, asymptotic properties on infinite intervals have been investigated by Griepentrog and M-arz [4]. Among other things, the notion of contractivity and that of B-stability were generalized to the case of DAE's and criteria for total stability were formulated. Algebraically stable IRK(DAE) were shown to be B-stable for index-1 DAE's, too, provided that the nullspace N of the leading Jacobian was constant. If this nullspace rotates, stability properties may change. In this paper, we study general linear index-2 DAE's exclusively, where the nullspace N := ker A(t) is assumed to be independent of t. A(t) and B(t) are assumed to be continuous in t. Equation (1.1) is not assumed to be in Hessenberg form and the coefficients A(t) and B(t) need not commute. Recall that Hessenberg index-2 DAE's have the special form This corresponds to the special coefficient matrices in (1.1) Moreover, it corresponds to a trivially constant nullspace N , since A(t) itself does not vary with t. This paper is a heavily revised and enlarged version of an earlier manuscript with the same title (Preprint 94-5). y Humboldt-Universit?t zu Berlin, Institut f?r Mathematik, D-10099 Berlin, Germany M. HANKE AND E. IZQUIERDO MACANA AND R. M - ARZ Presenting statements on the linear case we hope, as in the case of regular ordinary differential equations (ODE's), that it will be possible to carry over some properties to nonlinear DAE's via linearization. As far as we know, in case of index-2 DAE's stability analyses of integration methods on infinite intervals have only been presented for linear systems (see M-arz and Tischendorf [13], Wensch, Weiner, and Strehmel [16]). The latter paper is restricted to special Hessenberg form systems and relies on the so-called essentially underlying ODE introduced in Ascher and Petzold [2] for these special systems. We consider this case in Section 3 and describe the close relation between the inherent regular ODE which we will take up from [8] in Section 2, and the essentially underlying ODE in detail. Although the above mentioned paper [2] is not concerned with asymptotic stability on infinite intervals, it contains an observation that is highly interesting for us: Among other things, Ascher and Petzold point out that the backward Euler method applied to (1.2) may yield rather an explicit Euler formula for the essentially underlying ODE, and they discuss the influence of the blocks B on this phenomenon. We will show that not the derivatives of B 12 (t); B 21 (t), but the derivatives of the projector matrix H(t) := B 12 (t)(B 21 (t)B constitute the essential term, i.e., the rotation velocity of the subspace described by H(t) is the decisive feature. Our paper is aimed at explaining the importance of additional subspaces for answering questions concerning the asymptotic behaviour of integration methods. Hence, besides introducing the necessary fundamentals, Section 2 provides new results on the asymptotic stability of DAE solutions in the sense of Lyapunov as well as on contractive DAE's. In Section 4, BDF, IRK, and PIRK are investigated in detail. Asymptotic properties like A-stability and L-stability are shown to be preserved if a certain subspace it does not rotate. Moreover, we show that an algebraically stable IRK(DAE) is B-stable under these conditions. Section 5 illustrates our results by means of examples. For convenience of the reader, the short appendix provides the basic linear algebra facts once more. 2. Linear continuous coefficient index-2 equations. Consider the linear equation with continuous coefficients. Assume the nullspace of A(t) 2 L(R m ) to be independent of t and let Furthermore, set Obviously, S(t) is a subspace of R m which contains the solutions of the homogeneous form of the DAE (2.1). Note that the condition ASYMPTOTICS IN INDEX-2 DAE'S 3 characterizes the class of index-1 DAE's (see Appendix for related facts from linear algebra). Equation (2.2) implies that the matrix is nonsingular for all t 2 J , where Q projector onto N . Let Higher index DAE's are characterized by nontrivial intersections S(t) " N or equivalently by singular matrices G 1 (t). Definition. The DAE (2.1) is said to be index-2 tractable if the following two conditions hold, where In the following, let Q 1 (t) denote the projector onto N 1 (t) along S 1 (t), and (t). Due to the decomposition (2.5), Q 1 (t) is uniquely defined. Remarks. 1. It holds that dimN 1 2. Due to Lemma A.1, (2.4) and (2.5) imply that the matrix is nonsingular. But G 1 (t) is singular, independently of how Q is chosen [8]. 3. Applying Lemma A.1 once more we obtain the identities 4. Each DAE (2.1) having global index 2 is index-2 tractable with a continuously differentiable assuming Q 1 to belong to the class C 1 in the sequel is not restrictive. The conditions (2.4), (2.5) imply the decompositions which are relevant for the index-2 case. Taking this into account we decompose the DAE solution x into Multiplying (2.1) by PP respectively, and carrying out a few technical computations, we decouple the index-2 DAE into the system 4 M. HANKE AND E. IZQUIERDO MACANA AND R. M - ARZ Equation (2.10) represents the inherent regular ODE of the DAE system. On the other hand, if we consider (2.10) separately from its origin via the decomposition (2.9), we know that imPP 1 (t) is an invariant subspace of this explicit ODE in u. To be more precise: If we have at some t 0 2 J , then (2.10) implies Furthermore, (2.12) and (2.11) lead to respectively. Thus, solving (2.10) - (2.13) and setting we obtain the solutions of the DAE (2.1). Inspired by the above decoupling procedure, we state initial conditions for (2.1) as This yields but we do not expect x(t 0 Next, we shortly turn to the case of a homogeneous equation (2.1): For the system The matrix \Pi(t) := is also a projector, and is said to be the canonical projector for the index-2 case. Now, the following assertion is easily proved by means of the decoupling explained above. Theorem 2.1. Let (2.1) be index-2 tractable with continuously differentiable Q 1 . Then it holds: (i) The initial value problems (2.1), (2.14) are uniquely solvable in N provided that q 2 (ii) If x(:) solves the homogeneous equation, then it holds that (iii) Through each x exactly one solution of the homogeneous equation at time t 2 J . The solution space M (t) is a proper subspace of S(t) and Remarks. 1. The inherent regular ODE (2.10) is determined by the complete coefficient not only by its first term PP 2 B. If PP 1 (t) varies rapidly with t, the second term (PP 1 may be the dominant one. This should also be taken into account when considering the asymptotic behaviour of solutions of (2.1). ASYMPTOTICS IN INDEX-2 DAE'S 5 2. In general, the linear DAE (2.1) appears to be much simpler if the relevant subspaces N , N 1 , S 1 and the two projectors Q, Q 1 are constant. In that case (2.10) simplifies to 3. The value x 0 involved in the initial condition (2.14) is not expected to be a consistent initial value. What we have is As shown above, a consistent initial value for the homogeneous equation always belongs to M (t 0 ), which is precisely the set of consistent initial values then. 4. If the product PP 1 is time invariant, we have hence Note that (QP 1 G \Gamma1 2 B)(t) is also a projector onto ker A(t). It should be mentioned that the solution space M (t) remains time-invariant provided that both projectors are constant. Now we turn to the asymptotic behaviour of the solutions of the homogeneous equation. Considering the decoupled system (2.10) - (2.12) once more, we see that the component represents the dynamic one. Supposed the canonical projector remains bounded on the whole interval 1), the asymptotic behaviour of the solution is completely determined by that of its component u(t). Clearly, if u solves a constant coefficient regular ODE, we may characterize asymptotics by means of the corresponding eigenvalues. This is what we try to realize for the DAE in the following theorem. Theorem 2.2. Let (2.1) be index-2 tractable, Q 1 be of class C 1 , (PP 1 (i) Then the pencil -A(t) + B(t) has the eigenvalues - uniformly for implies each homogeneous equation solution to tend to zero as t !1, provided that the projector \Pi(t) remains uniformly bounded. Proof. Due to our assumptions, the inherent regular ODE has the constant coefficient 2 B. On the other hand, the nontrivial eigenvalues of \GammaP P (that is, eigenvalues that do not correspond to ker PP 1 ) are exactly the pencil eigenvalues of -A +B (cf. [10]). Let U (\Delta) denote the fundamental solution matrix of u Taking the solution representation into account, the assertion follows right away. Roughly speaking, the assumptions that PP 1 and PP have to be constant mean that there is a constant coefficient inherent regular ODE and a possible time dependence of the system may be caused by (time dependent) couplings only. 6 M. HANKE AND E. IZQUIERDO MACANA AND R. M - ARZ Next, what about contractivity in case of index-2 DAE's? In the regular ODE theory, contractivity is well-known to permit very attractive asymptotic properties of numerical integration methods. Corresponding results are obtained for index-1 DAE's in [4] by means of an appropriate contractivity notion. In particular, this notion says that a linear index-1 DAE (2.1) is contractive if there are a constant c ? 0 and a positive-definite matrix S such that the inequality holds true for all Here, we have used the scalar product hz; vi S := hSz; vi and the norm jzj S := hz; zi 1=2 S . Clearly, this reminds us of the one-sided Lipschitz condition used for contractivity in the regular ODE case (i.e. I in (2.1)). In the latter case we have, with However, things are more difficult for index-2 DAE's. First, considering the decoupled system again, we observe that each solution of the homogeneous DAE (2.1) satisfies the identities inspired by the notion of contractivity given for the index-1 case in [4], we state the following definition. Definition. The index-2 tractable DAE (2.1) is called contractive if the following holds: There is a constant c ? 0 and a symmetric positive-definite matrix S such that imply As usually, with this notion of contractivity, too, we aim at an inequality for all solutions of the homogeneous DAE, that shows the component decrease in that norm. The following theorem will show: If the canonical projector \Pi(t) is uniformly bounded, then the complete solution x(t) decreases. Theorem 2.3. Let (2.1) be index-2 tractable, Q 1 belong to C 1 , \Pi(t) be uniformly bounded on J and (2.1) be contractive. Then, it holds for each solution of the homogeneous equation that where fl is a bound of j\Pi(t)j S . ASYMPTOTICS IN INDEX-2 DAE'S 7 Proof. We have Not surprisingly, we obtain Corollary 2.4. Let (2.1) be index-2 tractable with continuously differentiable uniformly bounded \Pi(t). If the condition is satisfied for all u the estimate (2.21) is valid. Proof. It may be checked immediately that (2.19) and (2.22) lead to (2.20), i.e., implies contractivity. Note that there is no need for assuming (2.22) for all u . For the assertion of Corollary 2.4 to become true, it is sufficient that (2.22) holds for all u 2 im PP 1 (t), only. Inequality (2.22) looks like the usual contractivity condition for the regular ODE (2.10), i.e., the inherent regular ODE of (2.1). The only difference is that the values are taken from the subspace imPP 1 (t) instead of all of R m . Roughly speaking, one has: The DAE (2.1) is contractive if the inherent regular ODE (2.10) is contrctive on the subspace imPP 1 (t). As a direct consequence of other results on stability ([15], e.g.) one can deduce counterparts for linear index-2 DAE's, e.g. the well-known Poincar'e-Lyapunov Theorem 3. Specification of the projector framework for index-2 Hessenberg- form DAE's. Most authors restrict their interest to so-called Hessenberg-form equa- tions, i.e., to systems In our context this corresponds to I Obviously, z lar, which is the well-known Hessenberg-form index-2 condition. Under this condition the block is also a projector. It projects onto im B 12 (t) along ker B 21 (t). . It holds Furthermore, one has 8 M. HANKE AND E. IZQUIERDO MACANA AND R. M - ARZ The canonical projector \Pi is Recall that M precisely the solution space of the homogeneous form of (3.1). It is time dependent if the projector H(t) is. However, M (t) may also rotate with t even if H(t) is independent of time. Note that PP 1 is easier to compute than \Pi. Furthermore, the nontrivial part (i.e., dropping the zero rows) of the inherent regular ODE (2.10) reads now as where us emphasize once more that quickly varying subspaces may cause the term H 0 to dominate within this regular ODE. Clearly, H 0 u 1 corresponds to the term (PP 1 Theorems 2.1 and 2.2 apply immediately. In particular, we obtain: Suppose H(t) and are time-invariant. Then the eigenvalues of determine the asymptotic behaviour of the solution. us turn to the discussion of aspects of contractivity. For index-2 Hes- senberg-form DAE's (3.1), relation (2.20) applies to the first components only, i.e. should be satisfied if (cf. (2.19)) Moreover, (2.22) simplifies to for all Again we see that the constant-subspace case H 0 (t) j 0 becomes much easier. It should be stressed that the above decoupling as well as the inherent ODE are stated in the original coordinates. In particular, the subspace M (t) ae R m is precisely the one that contains the solutions of the original DAE. No coordinate transformation is applied and only a decomposition into characteristic components is employed. Ascher and Petzold [2] use a different approach to decouple characteristic parts of linear index-2 Hessenberg systems: They use a coordinate change z such that and ASYMPTOTICS IN INDEX-2 DAE'S 9 (cf. also [16]). In [2] the matrices R and S are constructed in the following way. Let First, a matrix R with linearly independent rows is chosen so that is satisfied. As a consequence, the m 1 \Theta m 1 block is nonsingular. Choosing S in such a way that hold true, we have The relation z 3 the main idea of that transformation. R(t) and S(t) are assumed to be smooth. Carrying out a few straightforward computations one obtains a regular ODE for the component z namely z 0 Equation (3.4) is said to be the essentially underlying ODE (EUODE) of the DAE (3.1). What does the EUODE have in common with the inherent regular ODE? What is the difference? Multiplying the EUODE (3.4) by S and taking into account that is given, we obtain (3.3). On the other hand, scaling the inherent regular ODE (3.3) by R leads back to the EUODE (3.4) because Thus, the EUODE turns out to be nothing else but a scaled version of the inherent regular ODE and vice versa. Due to the uniformly traced back to R m1 \Gammam 2 . Thus, the EUODE has the advantage to be written in the minimal coordinate space R m1 \Gammam 2 . Unfortunately, the matrices R and S are not uniquely determined. Consequently, the EUODE is strongly affected by the choice of R, S. Note that once an R is chosen, we may multiply by any regular K 2 L(R m1 \Gammam 2 ) to obtain another one by ~ R := KR. From this point of view, the inherent regular ODE (3.3) seems to be more natural, since all its terms are uniquely determined by the original data. is a direct component of the original variable x 1 , but the ODE (3.3) lives in the higher-dimensional space R m1 , and im(I \Gamma represents an invariant subspace. Ascher and Petzold [2] observed that the Euler backward method applied to the may behave like an explicit Euler method. Choose which simplifies the EUODE to z 0 (3. M. HANKE AND E. IZQUIERDO MACANA AND R. M - ARZ Via the transform z, the Euler backward formula applied to this special DAE If additionally B 12 (t) and R 0 (t) do not vary with t, (3.6) simplifies toh This is the explicit Euler formula for (3.5). Clearly this phenomenon is closely related to time-varying blocks R(t) and S(t) of the coordinate transformation. Let us mention again that this behaviour depends on the choice of R and S. In the following section we show that the behaviour of the characteristic subspace in the general case is decisive for understanding what really happens. 4. Asymptotic stability of integration methods. A number of widely used notions for the characterization of asymptotic properties of integration methods for explicit ODE's relies on the complex scalar test equation The asymptotic behaviour of a numerical method applied to (4.1) characterizes the asymptotics in the case of linear constant coefficient systems Here, the role of - is replaced by the eigenvalues of \GammaB. The justification for restricting the consideration to (4.2) is given by Lyapunov's theory: The linearization of a nonlinear autonomous explicit system at a stationary point provides criteria for the asymptotic behaviour of solutions. In essence, the same is true for index-1 and -2 DAE's [12]. Therefore, we are led to the constant coefficient DAE with regular matrix pencil -A +B. This equation can be transformed into the Kronecker canonical form z 0 I z 0: J 0 is a nilpotent matrix (J k discretization and transformation to (4.4) commute for many methods, the numerical solution for z vanishes identically, whereas y is discretized like an explicit system. Hence, numerical methods applied to constant coefficient linear DAE's trivially preserve their asymptotic stability properties that are based on the test equation (4.1) (e.g. A-, A(ff)-, L-stability). Thus, at first glance, one could expect the well-known concepts of asymptotics in the numerical integration of explicit ODE's to be sufficient for DAE's, too. How- ever, as described in Sections 2 and 3, DAE's have a more difficult structure than explicit ODE's, even in view of numerical integration. Roughly speaking, we should ASYMPTOTICS IN INDEX-2 DAE'S 11 not expect the numerical methods to match the subspace structure exactly if those subspaces rotate. The scalar test equation (4.1) turns out to be an inappropriate model in case of DAE's. Similar results about B-stability are more difficult to obtain. It is well-known that so-called algebraically stable Runge-Kutta methods are B-stable [6, p. 193] for explicit systems. In [4, p.129] a similar result is shown to be true for index-1 DAE's provided that (i) the nullspace N (A(t)) of A(t) does not depend on t, and (ii) the Runge-Kutta method is a so-called IRK(DAE) (a stiffly accurate method [6, p. 45]). There are simple linear examples showing that the backward Euler method loses its B-stability if (i) is not valid. We recall the notion of B-stability for DAE's having a constant leading nullspace: Definition [4]. The one-step method x called B-stable if for each contractive DAE the inequalities and jQx (1) are satisfied. Here, K ? 0 is a constant and x (1), x (2)are arbitrary consistent initial values. 4.1. BDF applied to linear index-2 DAE's. The k-step BDF applied to reads as At each step, equation (4.5) provides an approximation x ' of the exact solution value Recall that the nullspace of A(t) is assumed to be constant. Supposed (2.1) is index-2 tractable, we may decouple (4.5) and (2.1) simultaneously (cf. ARZ where we have used the above decomposition again, i.e., In particular, if the inhomogeneity q vanishes identically, then the Q 1 -components are both zero, and one hash for approximation of and for approximation of The following proposition is an immediate consequence. Proposition 4.1. Let (2.1) be index-2 tractable with continuously differentiable . Then the BDF method applied to (2.1) generates exactly the same BDF method applied to the inherent regular ODE (4.11) if and only if the projector PP 1 (t) does not vary with t. For a constant projector PP 1 , the BDF methods retain their asymptotic stability properties for index-2 DAE's provided the canonical projector \Pi(t) remains uniformly bounded. On the other hand, varying subspaces may cause the term PP 0 1 to dominate the inherent regular ODE itself. For instance, the backward Euler method provides thenh which shows that u(t ' may or may not happen. As it was mentioned in Section 3, Ascher and Petzold [2] have observed this phenomenon in case of linear index-2 Hessenberg systems (3.1) (cf. also Section 3). However, this is not surprising since we cannot expect any discretization method to follow the subspaces precisely without profound information on the inner structure of the DAE. Naturally, similar arguments apply to Runge-Kutta methods, too. 4.2. Implicit Runge-Kutta methods and their projected counterparts applied to linear index-2 DAE's. According to the originally conceived method for the numerical solution of ordinary differential equations, an implicit Runge-Kutta (IRK) method can be realized for the DAE (2.1) in the following way [14]: Given an approximation x '\Gamma1 of the solution of (2.1) at t '\Gamma1 , a new approximation x ' at obtained via s ASYMPTOTICS IN INDEX-2 DAE'S 13 'i is defined by and the internal stages are given by s s: The coefficients a ij , b i , c i determine the IRK method, and s represents the number of stages. Assume the matrix A := (a ij i;j=1 to be nonsingular and denote its inverse s s Equations (4.12) - (4.14) are equivalent to s s s s: Looking at (4.16) we observe that the internal stages do not depend on Qx '\Gamma1 . The special class of IRK methods (IRK(DAE)) with coefficients is shown to stand out from all IRK methods in view of their applicability to DAE's in that case, the new value x belongs to the obvious constraint manifold Therefore we have For Hessenberg equations (3.1), relation (4.18) simplifies to In general, if (4.17) is not fulfilled, then we have % 6= 0, and (4.18) resp. (4.19) are no longer true. Since this behaviour is a source of instability (for h ! 0), Ascher and Petzold [1] propose another version for the application of IRK methods to index-2 Hessenberg DAE's (2.18), the so-called Projected IRK (PIRK). Actually, after realizing the standard internal stage computation, the recursion (4.15) is now replaced by s s and - ' is determined by 14 M. HANKE AND E. IZQUIERDO MACANA AND R. M - ARZ If we multiply (4.20) by I \Gamma H(t ' ), - ' can be eliminated: s s On the other hand, (4.21) is equivalent to It should be mentioned that for IRK(DAE) the projected version is exactly the same as the original one, since (4.17) implies - Considering (4.22) - (4.23) in association with the projector formulae (3.2), an immediate generalization of PIRK methods to fully implicit linear index-2 systems (2.1) is suggested by s s Since the internal stages - X 'j do not depend on Q-x '\Gamma1 , there is no need to compute Q-x ' at this stage. Now return to the standard IRK (4.15) - (4.16) and decouple (4.16) in the same way as (2.12). For that, we decompose A straightforward computation yieldsh s s \GammaP s s (4.26)h s s s The recursion (4.15) can be decomposed simply by multiplying by the projections: s s ASYMPTOTICS IN INDEX-2 DAE'S 15 s s s s s s s s Now, consider the homogeneous case, that is we set If the inhomogeneity q vanish identically, then v does so, too. Moreover, all values V 'i are equal to zero. However, if % 6= 0, this is no longer true for This means that, in general, the resulting x ' has a nontrivial component in contrast to the exact solution that fulfills Q In more detail, (4.26) reduces toh s s \GammaP s which supposedly approximates Moreover, (4.27) yields s \GammaQ s for approximating In the consequence, the following result holds true for IRK methods analogously to Proposition 4.1 for the case of BDF methods: Proposition 4.2. Let (2.1) be index-2 tractable with continuously differentiable . Then the IRK method applied to (2.1) generates exactly the same IRK method applied to the inherent regular ODE (4.30) if and only if PP 1 (t) does not vary with t. For constant PP 1 , the solution (4. M. HANKE AND E. IZQUIERDO MACANA AND R. M - ARZ of the homogeneous equation is approximated at t ' by s s (4.39)h s s s Starting with a consistent initial value x 0 (with the components v ' vanish step by step, too. For IRK(DAE), (4.42) provides that is, in the case of constant PP 1 , the approximation x ' belongs to the solution manifold M (t ' ) given in Theorem 2.1. Let us briefly turn to PIRK methods (4.24), (4.25). For homogeneous equations, The decoupled system parts (4.32), (4.34) remain valid also for the "-" values. Proposition 4.3. Proposition 4.2 is true for PIRK methods, too. It should be mentioned that, for constant PP 1 , in PIRK methods we have simply instead of (4.41). Ascher and Petzold [1] have not considered a recursion for the component Q-x ' for Hessenberg systems (2.19). Nevertheless, if one is interested in approximations Q-x ' , a recursion like (4.42) will come up again. In that case, the only difference between PIRK and IRK methods is the determination of the Q 1 -components versus (4.41)). Note again that PIRK and IRK are identical for IRK(DAE). Next, concerning B-stability, the following assertion shows the notion of contractivity given in Section 2 to be useful. Theorem 4.4. Let (2.1) be index-2 tractable with continuously differentiable Q 1 , each algebraically stable IRK(DAE) applied to (2.1) is B-stable. Proof. Denote . Due to the algebraical stability, positively semi-definite matrix. Since we deal with linear DAE's only, it remains to show the inequalities jP x ' for the case of the homogeneous equation (2.1). ASYMPTOTICS IN INDEX-2 DAE'S 17 therefore Additionally, with an IRK (DAE) we also have s s since holds true for all ~ t. Hence, using the contractivity (cf. Section 2) we obtain the inequalities s: Now, following the standard lines, we compute s s 'i s s s 'i s s 'i s s Finally, x It should be noted that Theorem 4.4 does not apply to PIRK. While the first part, i.e., jP x ' holds true analogously, the necessary relation for the nullspace component is not given at all for % 6= 0. 5. A numerical counterexample. In the previous sections we have seen that BDF and Runge-Kutta methods preserve their stability behaviour if PP 1 is constant. The following example shows that these properties get lost if PP 1 varies with time. Consider the DAE with M. HANKE AND E. IZQUIERDO MACANA AND R. M - ARZ where are constant. Note that (5.1) is an index-2 Hessenberg system. One easily computes (using such that Compute the projections Taking into account that in (5.1), the inherent regular ODE (2.10) reads The solution subspace M (t) (cf. Theorem 2.2) is given by subject to consistent initial values (2.13) may be reduced to the scalar ODE, together with Consequently, the asymptotic stability of (5.2) is governed by the sign of - (indepen- dently of j 2 R). The parameter j measures the change of N 1 (t). fi serves only for mixing the P component with the nullspace component. Now the complete solution of (5.1) can be easily computed using (2.10) - (2.12). If x 0 2 R 3 is a consistent initial value at the solution of (5.1) is was solved using the 5-step BDF (Fig. 5.1) and an algebraically stable 2-stage Runge-Kutta method introduced by Crouzeix (cf. [5, p. 207]) with ae - \Gamma0:73 5.2). The figures show the norm of the numerical solution at the end of the different values of j and -. Note that, for a constant coefficient system. The results indicate that the asymptotic behaviour of the numerical solution depends not only on the asymptotic stability of the differential equation (5.3) (controlled by -), but also on the geometry of the problem (controlled by j). ASYMPTOTICS IN INDEX-2 DAE'S 19 Appendix Basic linear algebra lemma. A basic connection between the spaces appearing in the tractability index and the choice of the corresponding projectors is given by the following lemma, which may be directly obtained from Theorem A.13. and Lemma A.14. in [4]. Lemma A.1. Let - be a projector onto ker( - A). Denote - A)g. Then the following conditions are equivalent: (i) The matrix - Q is nonsingular. A). If - G is nonsingular, then the relation holds for the canonical projector - along - S). Proof. (i) ! (ii) The space R m can be described as - A), because holds for any z 2 R m . z 2 obviously lies in ker( - Q is a projector onto A). For z 1 we obtain S. It remains to show that - f0g. To this end, let x 2 - A). Qx holds and there exists a z 2 R m such that - Qx and Qx. Consequently, This holds trivially by definition. chosen such that - Ax and so S. On the other hand, - Qx lies in ker( - A). Thus, x 2 holds due to the assumption. That means, - Q). Then has to be true, and G is nonsingular. Because of the uniqueness of the decomposition ( ), the latter assertion follows immediately. --R Projected implicit Runge-Kutta methods for differential-algebraic equations Stability of computational methods for constrained dynamic systems Numerical Solution of Initial-Value Problems in Differential-Algebraic Equations Solving Ordinary Differential Equations I Solving Ordinary Differential Equations II Approximation von Algebro-Differentialgleichungen mit bereich Mathematik Order results for implicit Runge-Kutta methods applied to differential algebraic systems Nonlinear Differential Equations and Dynamical Systems Stability investigations for index-2-systems --TR --CTR Roswitha Mrz , Antonio R. Rodrguez-Santiesteban, Analyzing the stability behaviour of solutions and their approximations in case of index-2 differential-algebraic systems, Mathematics of Computation, v.71 n.238, p.605-632, April 2002 Hong Liu , Yongzhong Song, Stability of numerical methods for solving linear index-3 DAEs, Applied Mathematics and Computation, v.134 n.1, p.35-50, 10 January I. Higueras , R. Mrz , C. Tischendorf, Stability preserving integration of index-1 DAEs, Applied Numerical Mathematics, v.45 n.2-3, p.175-200, May I. Higueras , R. Mrz , C. Tischendorf, Stability preserving integration of index-2 DAEs, Applied Numerical Mathematics, v.45 n.2-3, p.201-229, May Roswitha Mrz, Differential algebraic systems anew, Applied Numerical Mathematics, v.42 n.1, p.315-335, August 2002

asymptotic properties;backward differentiation formulas;differential-algebraic equation;runge-kutta method;stability

296420

On Improving the Convergence of Radau IIA Methods Applied to Index 2 DAEs.

This paper presents a simple new technique to improve the order behavior of Runge--Kutta methods when applied to index 2 differential-algebraic equations. It is then shown how this can be incorporated into a more efficient version of the code {\sc radau5} developed by E. Hairer and G. Wanner.

Introduction In recent years, differential algebraic equations (DAEs) have been studied by various authors (see [HW91], [HLR89], [BCP91]), and their importance acknowledged by the development of specific solvers such as DASSL from [Pet86] or RADAU5 from [HW91]. An especially important class of DAEs arising in practice are semi-explicit systems of the form where g y f z is assumed to be of bounded inverse in a neighbourhood of the solution of (S). Here, we are interested in obtaining a numerical approximation to (S) accurate for both the differential and the algebraic components. Although some of the ideas presented in this paper also apply to more general Runge-Kutta methods, we will focus on Radau IIA methods, given that they were used to build the code RADAU5. Their construction as well as some of their properties are briefly recalled in Section 1.1. When applying a s-stage Radau IIA method to (S), the orders of convergence are respectively for the y-component and s for the z-component (see [HLR89]). In some situations, where getting an accurate value of z may be important (in mechanics for instance), one is led to use a different approach. Generally speaking, the order reduction phenomenon may be overcome by the following techniques: (i) a first possibility consists in applying the Radau IIA method to the index one formulation (S) of Since Radau IIA methods applied to index one DAEs exhibit full order of convergence for y and z (see Theorem 3-1 [HLR89]), the order of convergence is now 2s \Gamma 1 also for z. However, solving (S) can be considerably more costly: as a matter of fact, this requires to evaluate the Jacobian of the function F (y; at each step (or whenever the convergence rate of the Newton iteration gets too small), instead of the function F (y; Another drawback is that the numerical solution is not forced any longer to lie on the constraint manifold (ii) a second idea consists in computing the z-component by solving the additional equation g y (y)f(y; z), i.e. in projecting the numerical solution on the so-called "hidden constraint". The corresponding numerical scheme now reads The order of convergence for the y-component is still 2s \Gamma 1 and can be shown to be now also for the z-component by using the Implicit Function Theorem. However, this technique is once again computationally more demanding than the original one: solving the new implicit part of (S n ) requires an accurate evaluation of g y (y) at each step, and not only, as previously mentioned, whenever the convergence rate of the Newton iteration becomes too small. It can be nevertheless noted that in a parallel environment, this additional cost would be shadowed by the use of a second processor. In this paper, we present a third approach which does not require an analytical form of g y and whose computational cost is basically equal to what it is for the standard formulation. It is based on the PI n-976 4 A. Aubry , P. Chartier observation that the errors in the z-component are essentially of a local nature, at least up to the order of convergence of the y-component. As a consequence, making z more accurate is a matter of recovering the significant terms that appear in the so-called "B-series of the error". This is made possible by considering the composition of the basic method with itself over several steps. It can be noted that similar ideas were used by R.P.K. Chan to deal with the order reduction of Gauss methods when applied to certain stiff problems (see [BC93]). The new order conditions will be determined in Section 2. While they can be derived in a straightforward manner from the work of E. Hairer, C. Lubich and M. Roche ([HLR89]), they are still relatively unknown since they are not satisfied by most classical Runge-Kutta methods (see also [BC93] or [CC95]). It will then be shown that some of those conditions are actually redundant and may be omitted. This is a crucial aspect of the method, since it allows the construction of formulas with a manageable level of complexity. In Section 3, the implied modifications to the code RADAU5 are listed. Finally, numerical results are presented that illustrate the advantages of this new technique. 1.1 Basic properties of Radau IIA methods Radau IIA methods can be defined by the s A particular method R will be characterized by the triple (A; b; c) where (a ij ) i;j=1;\Delta\Delta\Delta;s is a s \Theta s matrix, s-dimensional vector and s-dimensional vector. In the sequel, we will furthermore use the notations 1.1.1 Construction of Radau IIA methods Their coefficients are uniquely determined by the following conditions: 1. are the ordered zeros of the Radau right polynomial dx x 2. 3. The coefficients a ij of the matrix A satisfy C(s): As the c i are all distinct, A is non-singular. 1.1.2 Some useful properties We will refer here to the additional simplifying assumptions D(-) introduced, as B(p) and C(j) of previous subsection, by J.C. Butcher (see [HW91] page 75). Irisa On improving the convergence of Radau IIA methods applied to index 2 DAEs 5 1. Due to the conditions C(s), B(s) and c methods are stiffly accurate (i.e. vectors b and (a are solution of the same Cramer system). 2. They are collocation methods (Theorem 7-8 [HNW91], page 212). 3. 4. D(s \Gamma 1) is satisfied (Lemma 5-4 [HW91], page 78). Optimal convergence results have been obtained for those methods by J.C. Butcher on the one hand (Theorem 5-3 of [HW91]) and E. Hairer, C. Lubich and M. Roche on the other hand (Theorems 3-1, 4-4 and 4-6 of [HLR89]). They are collected in Table 1. Table 1: Optimal global error estimates for the s-stage method Radau IIA 1.2 Increasing the order of convergence of the z-component When applying a s-stage Radau IIA method c) to the system (S), we obtain z are the coefficients of the matrix A \Gamma1 . In (4), z n vanishes because us now replace the vector (b T A by an adjustable vector (4). By doing so, we define a new method Rw . It is easily seen that the order of convergence of the y-component remains unchanged. As for the z-component, the lack of accumulation makes the errors purely local. The convergence behaviour of the z-component is thus entirely determined by the following order conditions from Theorem 8-6 and 8-8 of [HW91]. Proposition 1 Let ffi y and ffi z be the local errors respectively for the y and z components of a Runge-Kutta method. Then we have, where DAT2 y , DAT2 z are sets of trees, \Phi a vectorial function and fl; ae scalar ones associated with the trees 1 . Nevertheless, w has not enough components to allow for an order of convergence greater than s is the optimal vector). Hence, to get sufficient freedom, we need to consider the composition R oe of R over oe steps. As variable steps h i , are considered, we also have to introduce the ratios r characterized by the triple (A; B; C) where ffl A is the blockmatrix blocks of the form 1 See [HW91] for a definition of these notions. PI n-976 6 A. Aubry , P. Chartier ffl B is the blockvector (B i ) i=1;\Delta\Delta\Delta;oe with blocks of the form B ffl C is the blockvector (C i ) i=1;\Delta\Delta\Delta;oe with blocks of the form C e. Replacing by an adjustable vector w offers oe \Theta s degrees of freedom, i.e hopefully enough for oe ? 1 to increase the order of convergence of the z-component. It is our aim now to show how to construct w and how to implement the new method R oe;w . 2 Construction of the vector w In this section, the effective construction of w is described. It should be emphasized that its components depend on r forcing one evaluation per step. However, these additional computations become negligible as soon as the dimension of the system (S) is large enough. 2.1 Order conditions The conditions for order k are enumerated below together with the associated trees. required with z - are required with s z - z - Let U s be AC s+1 C s+1 . If (s) is satisfied, then (C s+1 ) is equivalent to are required with 2 z - z - s z - z - s z - Let U s+1 be AC equivalent to 2 The dot stands for the componentwise product. Irisa On improving the convergence of Radau IIA methods applied to index 2 DAEs 7 are required with z - z - z - z - z - s z - z - s z - s z - z - s z - s z - z - Let U s+2 be AC s+3 C s+3 . If (s + 7) is satisfied, then (C s+3 ) is equivalent to These conditions are obtained by Proposition 1 and by using simplifying assumptions (it is important to note that the composite method R oe satisfies B(2s \Gamma 1), C(s) and D(s \Gamma 1)). 2.2 Preliminary calculus In order to later simplify the equations for w, we now state some basic results. eb T , then F Proof: By definition, we have eb T A \Gamma1 . As the method c) is stiffly accurate, ve T s where ve T eb ve T ve T ve T ve T be the s \Theta s blocks of the matrix A \Gamma1 , then PI n-976 8 A. Aubry , P. Chartier Proof: By definition, AA Lemma 2 is then easily proved by induction on i. 2 oe;n be the oe \Theta s dimensional vector AC n+1 C n+1 and u n be the s dimensional vector is given by k=s r k+1 Proof: Let n be less than or equal to 2s \Gamma 2. For all S i;n can be expanded as follows r k+1 r k+1 e: Owing to C(s), we have k=s r k+1 e: Similarly, T i;n can be expanded as l r l implies that and as n - 2s \Gamma 2, we obtain r l+1 e e e On improving the convergence of Radau IIA methods applied to index 2 DAEs 9 Lemma 4 For all n less than Proof: This follows at once from B(2s \Gamma 1). 2 Lemma 5 Let oe;n be the vector A \Gamma1 U n , then for all integer n - by k=s Proof: By definition, X A \Gamma1 U i;n . Using Lemma 3, we obtain k=s r k+1 k=s and we use Lemma 4 to complete the proof. 2 Lemma 6 For all n less than 2s Proof: This follows straightforwardly from the order conditions for the trees z - z - 1. Let oe;n be the vector C:A k=s r k+1 2. Let oe;n be the vector A \Gamma1 (C:U n ); k=s r k+1 Proof: By definition, Y and the first part of the result is obtained by applying Lemma 5. Now, let oe;n denote the vector C:U n . From Lemma 3, we can write T i;n as k=s r k+2 We have furthermore Z so that Lemma 2 leads to k=s s n\Gammak s n+1\Gammak k=s r k+1 PI n-976 A. Aubry , P. Chartier The result then becomes a consequence of lemmas 4 and 6. 2 Lemma 8 For all n less than 2s Proof: This follows from the order conditions for the trees z - z - z - Proof: The results follow from the order conditions for the trees [[-; [ z - z - z - z - z - 2.3 Results for the 2-stage method In order to get a third-order method, w has to satisfy the following linear system (S L;2 Taking leads to a system with 4 equations and 4 unknowns. For convenience, we recall below the coefficients of the 2-stage Radau IIA method, R =312 The matrix M corresponding to (S L;2 ) is then of the form and we have Hence, for all r 1 2 (0; 1), M is non-singular and (S L;2 ) has the following unique solution Irisa On improving the convergence of Radau IIA methods applied to index 2 DAEs 11 2.4 Results for the 3-stage method In order to get a fifth-order method, w must satisfy the following linear system (S L;3 Taking now leads to a system with ten equations but only nine unknowns. However, we will show in this section, that one of these equations is identically satisfied. Collecting all results of Section 2.2, we get r 4 r 4 r 4 3 A r 4 r 4 r 4 r 4 r 4 r 4 r 4 r 4 r 4 It is found that r 4 r 4 r 4 r 4 r 4 r 4 and the system (S L;3 ) is equivalent to Theorem are linearly dependent. Proof: It is enough to show that v 1 , v 2 and u 3 are linearly dependent. Since c) is of order 5 for the differential component, we have so that b T u implies the result.2 the vectors A \Gamma1 U 3 , U 3 and V 1 (for example) become linearly dependent. To prove this, it is sufficient to show that A \Gamma1 u 3 , u 3 and v 1 are dependent. As b T A \Gamma1 u PI n-976 A. Aubry , P. Chartier 0, we can conclude as in Theorem 1. In this case, w depends on one parameter that can be chosen so as to minimize the quantity where DAT2 z 5g. This seems a natural goal to achieve, since this is an attempt to minimize the local error. For convenience, Table 2 collects the trees of DAT2 z (5) and the values of the associated functions ff, fl and \Phi. To compute the ff's, we refer to [Hig93]. tree u ff(u) fl(u) \Phi(u) Table 2: Trees of DAT2 z (5) and their associated functions 2.5 Sketch of the case To achieve order 7 for the algebraic component, we have to solve the system (S L;4 ) composed of (C 4 ), Section 2.1), that is to say 24 equations. Comparing the number of equations and the number of unknown, we could consequently think of taking In fact, it can be shown that is sufficient, since 5 equations are identically satisfied (see Section 6.1). However, it does not seem reasonable any more to consider a practical implementation of the corresponding method, owing to the complexity of the formulas for variable stepsize. Remark 2 If the stepsize is constant (i.e. r equations are identically satisfied (see Section 6.1), and oe can be chosen equal to 4. 3 Modifications to the code RADAU5 The 3-stage Radau IIA method has been implemented by E. Hairer and G. Wanner in order to solve problems of the form MY and DAEs of index less than or equal to three can be solved by this code, called RADAU5. A precise description is given in Section IV-8 [HW91] and we will adopt the notations used there. Implementing our method requires slight modifications to the subroutines "radcor" and "solout" which are actually replaced respectively by "radcorz" and "soloutz". Irisa On improving the convergence of Radau IIA methods applied to index 2 DAEs 13 3.1 Modifications to radcor Only the computation of the algebraic component (z) is modified (if an index 2 DAE is solved). Once the n th step has been accepted, two cases are considered: 1. less than three steps have been computed. Then, we keep the internal stages z n;1 , z n;2 , z n;3 and the step size h n . The value of (y 2. three or more steps have been computed. Then, we keep the internal stages z n;1 , z n;2 , z n;3 and the step size h n . r 1 , r 2 are computed by the formulas and r and w by the subroutine "vectw2". For (y n+1 ; z n+1 ) we put z Remark 3 For continuous outputs, we need also to keep the internal stages over two steps for the differential components (see section 3.2.2). 3.2 New subroutines 3.2.1 Vectw2 In order to compute the formal expression of w and to create the associated fortran subroutine, the manipulation package Maple was used. A call to vectw2 uses the format vectw2(icas,vw,r1,r2), where the inputs are one of the five cases described below (icas) and the parameters r 1 , r 2 (r1,r2). The output is the vector w (vw). Five cases are considered: 1. 1. In this case, we have seen that w depends on one parameter. It is optimized as explained in Remark 1. 2. r 3. r 4. r 5. r 1 This allows us to reduce the cost of computation and to eliminate computational problems: had we used the general expression of w (case 5), divisions by zero would have occured in the cases 1 to 4. 3.2.2 Continuous outputs In the code Radau5, the subroutine "solout" provides the user with approximations at equidistant output-points. The corresponding interpolation formulas are implemented in the subroutine "contr5". "contr5(I,x)" gives an approximation U I (X) to the I th component of the solution Y at the point x should lie in the interval [x n ; x n+1 ]). U is the collocation polynomial: it is of degree 3 and defined by 3: PI n-976 14 A. Aubry , P. Chartier For index 2 DAEs, are polynomials of degree 3 which satisfy 3: By Theorem 7-8 ([HW91]), we have As our aim is to increase the order of convergence for the algebraic component, it seems natural to search for an approximation P Approximation of z(x) Let x be of the form x n\Gamma2 We define q as follows where the vector satisfies the linear system Using the notations of section 2.4, we have Proposition 2 For all ' 2 (0; 1], S z L;3 (') possesses a solution and Proof: From Theorem 8-5 and 8-6 in [HW91], we have According to the analysis of section 2.1, this leads to the system S z L;3 (') which, by Theorem 1, possesses a solution. 2 The subroutine "vectwz" computes w('). As for the vector w, five cases are considered. When h depends on one parameter (case 1) which is not adjusted as in Remark 1. In this case, the value defined by continuity for w(') is choosen. A call to "vectwz" uses the format vectwz(icas,vwz,r1,r2,t), where the inputs are one of the five cases described before (icas) and the parameters r 1 , r 2 , ' (respectively r1,r2,t). The output is the vector w(') (vwz). Approximation of y(x) Let x be of the form x Irisa On improving the convergence of Radau IIA methods applied to index 2 DAEs 15 where the vector satisfies the linear system C) is the Runge-Kutta method R 2 obtained by the composition of the 3-stage Radau IIA method c) over two steps, re and E is the vector (z - Proposition 3 Proof: Let us introduce the vector and the following polynomial u: From Theorem 8-5 and 8-6 in [HW91], we have Using the points of the internal stages, it follows so that (P ) is equivalent to the system S y The subroutine "vectwy" computes B(j). A call to "vectwy" uses the format vectwy(vwy,r,t), where the inputs are the parameters r and j (respectively r,t) and the output the vector B(j) (vwy). The subroutine "soloutz" provides the user with approximations (p(x i out out out out at equidistant output-points out ) i=1;\Delta\Delta\Delta;N . For the differential component, x i out is of the form x choosen so as to satisfy x out - x n+1 and for the algebraic one, x i out is of the form x n\Gamma2 is choosen in satisfy x out - x n . A call to "soloutz" uses the format where the inputs are the number of accepted steps (nr), x n (xold), x n+1 (x), (y n+1 ; z n+1 ) (y), the system's dimension (neq), one of the five cases described before for the computation of w(') (icas), the parameters r1, r2 (r1,r2), the stepsize h n\Gamma2 variable (last) to indicate if the last computational step has been reached. PI n-976 A. Aubry , P. Chartier significant digits modified radau5 Figure 1: Precision versus computing time - algebraic components - test problem 4 Numerical experiments 4.1 Test problem We consider the index two problem: where \Psi is the following infinitely smooth function: with consistant initial values. The exact solution is In both codes, we set (work(4) is the parameter - in the stopping criterion for Newton's method. work(4), work(5) are the parameters c 1 , c 2 in the stepsize control, see [HW91], page 130-134). In figure 1, we plot the cpu time against the number of significant digits (\Gammalog 10 (absolute error)) of the algebraic components, for both codes. For this, we use continuous outputs : outputs are required at we compute the global error and then take the maximum over all values. In figure 2, we plot the cpu time against the number of significant figures of the differential components, for both codes. In the following problems, only the modified parameters will be shown. Irisa On improving the convergence of Radau IIA methods applied to index 2 DAEs 17261014 significant digits modified radau5 Figure 2: Precision versus computing time - differential components - test problem 4.2 Pendulum The simplest constrained mechanical system is the pendulum, whose equations of motions are described in [HW91], page 483-485. We have applied the code Radau5 and the modified code to the GGL (Gear, Gupta and Leimkuhler) formulation with consistant initial values simplicity, we took In figure 3 (respectively 4), we plot the cpu time against the number of significant digits of the algebraic (resp. differential) components, for both codes. 4.3 Multibody mechanism A seven body mechanism is described in [HW91], page 531-545. We have applied the code Radau5 and the modified code to the index 2 formulation with consistant initial values. In figure 5 (respectively 6), we plot the cpu time against the number of significant digits of the algebraic differential) components, for both codes. Here, outputs are required at t 4.4 Discharge pressure control This simplified model of a dynamic simulation problem in petrochimical engineering is described in [HLR89], page 116-118. We have applied the code Radau5 and the modified code to the following PI n-976 A. Aubry , P. Chartier246810 significant digits modified radau5 Figure 3: Precision versus computing time - algebraic components - significant digits modified radau5 Figure 4: Precision versus computing time - differential components - pendulum Irisa On improving the convergence of Radau IIA methods applied to index 2 DAEs 19 significant digits modified radau5 Figure 5: Precision versus computing time - algebraic components - seven body mechanism significant digits modified radau5 Figure versus computing time - differential components - seven body mechanism PI n-976 A. Aubry , P. Chartier13579 significant digits modified radau5 Figure 7: Precision versus computing time - algebraic components - discharge pressure control26100.01 significant digits modified radau5 Figure 8: Precision versus computing time - differential components - discharge pressure control with consistant initial values. Here, initial step size h is equal to 10 \Gamma5 . In figure 7 (respectively 8), we plot the cpu time against the number of significant digits of the algebraic (resp. differential) components, for both codes. Here, outputs are required at t Irisa On improving the convergence of Radau IIA methods applied to index 2 DAEs 21 5 Conclusion A new simple technique to overcome the order reduction phenomenon, appearing for the algebraic component when Radau IIA methods are applied to index two DAEs, is proposed. Increasing the order of convergence of z is made possible by considering the composition of the basic method with itself over oe steps. As z n+1 is defined in the basic method as a linear combination of the internal stages, a good choice of oe should provide enough freedom for the order conditions associated with the composite method to be satisfied. We have determined these order conditions which derive straightforwardly from the work of E. Hairer, C. Lubich and M. Roche (section 2.1). Then we have shown that some of those conditions are redundant and might be omitted for s-stage Radau IIA methods with s - 4 (section 2.3 to 2.5). It could be interesting to generalize this simplifications to any Radau IIA methods. A general question will be to determine how many compositions of a s-stage Radau IIA method have to be considered to obtain an order of convergence equal to 2s \Gamma 1 for the algebraic component. However, it does not seem reasonable any more to consider a practical implementation of the corresponding method for s - 4, owing to the complexity of the formulas for variable stepsize. The formulas for incorporated in the code Radau5 developed by E. Hairer and G. Wanner. Slight modifications were required (section 3). Only the computation of the algebraic component was modified and a new procedure in order to have continuous outputs was created where we have used our technique for both components (algebraic and differential) to compute approximations of order five at equidistant output-points. According to our numerical experiments, results for the differential components are disappointing. However, the use of our technique in the code Radau5 leads to an increase of the accuracy for the algebraic components (when tolerances are sufficiently small). 6 Appendix 6.1 Construction of the vector w in the case In this section, we explained the calculus of linear algebra used to show that composed five times the s-stage Radau IIA method is sufficient in the case expand the following vectors of (C 6 ) (We use Lemma A A A and the following vectors of (C 7 ) (We use Lemma PI n-976 22 A. Aubry , P. Chartier A r 6 r 6 A A A A r 6 r 6 r 6 r 6 Let introduce the following vectors (idem for the vectors t r 6 r 6 r 6 r 6 r 6 r 6 r 6 and the system (S L;4 ) is equivalent to Irisa On improving the convergence of Radau IIA methods applied to index 2 DAEs 23 Proposition 4 1. are linearly dependent. 2. are linearly dependent. 3. are linearly dependent. 4. are linearly dependent. 5. are linearly dependent. Proof: Because of the expression of the vectors it is sufficient to show that are linearly dependent (idem for the part 2 to 5 of the proposition). The method c) is of local order 8 for the differential component. Thus, order conditions associated with the trees of DAT2 y (7) are satisfied. In particular, Finally, we obtain but b 6= 0, hence the proposition is shown. 2 Remark 4 If the step size is constant (i.e. r 1. are linearly dependent. 2. are linearly dependent. 3. are linearly dependent. 4. are linearly dependent. Hence, nine equations are identically satisfied and p can be choosen equal to four. PI n-976 A. Aubry , P. Chartier --R On smoothing and order reduction effects for implicit Runge-Kutta formulae Numerical solution of initial value problems in differential-algebraic equations A Composition Law for Runge-Kutta Methods Applied to Index-2 Differential-Algebraic Equations Coefficients of the Taylor expansion for the solution of differential-algebraic systems The Numerical Solution of Differential Algebraic Systems by Runge-Kutta Methods Solving Ordinary Differential Equations (Vol Stiff Problems and Differential Algebraic Problems (Vol. A description of DASSL: A differential/Algebraic System Solver. --TR --CTR Frank Cameron , Mikko Palmroth , Robert Pich, Quasi stage order conditions for SDIRK methods, Applied Numerical Mathematics, v.42 n.1, p.61-75, August 2002 Frank Cameron, A Matlab package for automatically generating Runge-Kutta trees, order conditions, and truncation error coefficients, ACM Transactions on Mathematical Software (TOMS), v.32 n.2, p.274-298, June 2006

rooted trees;runge-kutta methods;composition;differential-algebraic systems of index 2;Radau IIA methods;simplifying assumptions

296955

Rapid Concept Learning for Mobile Robots.

Concept learning in robotics is an extremely challenging problem: sensory data is often high-dimensional, and noisy due to specularities and other irregularities. In this paper, we investigate two general strategies to speed up learning, based on spatial decomposition of the sensory representation, and simultaneous learning of multiple classes using a shared structure. We study two concept learning scenarios: a hallway navigation problem, where the robot has to induce features such as opening or wall. The second task is recycling, where the robot has to learn to recognize objects, such as a trash can. We use a common underlying function approximator in both studies in the form of a feedforward neural network, with several hundred input units and multiple output units. Despite the high degree of freedom afforded by such an approximator, we show the two strategies provide sufficient bias to achieve rapid learning. We provide detailed experimental studies on an actual mobile robot called PAVLOV to illustrate the effectiveness of this approach.

Introduction Programming mobile robots to successfully operate in unstructured environments, including offices and homes, is tedious and difficult. Easing this programming burden seems necessary to realize many of the possible applications of mobile robot technology [7]. One promising avenue towards smarter and easier-to-program robots is to equip them with the ability to learn new concepts and behaviors. In partic- ular, robots that have the capability of learning concepts could be programmed or instructed more readily than their non-learning counterparts. For example, a robot that could be trained to recognize landmarks, such as "doors" and "intersec- tions", would enable a more flexible navigation system. Similarly, a recycling robot, which could be trained to find objects such as "trash cans" or "soda cans", could be adapted to new circumstances much more easily than non-learning robots (for example, new objects or containers could be easily accommodated by additional training). Robot learning is currently an active area of research (e.g. see [5], [6], [9], [16]). Many different approaches to this problem are being investigated, ranging from supervised learning of concepts and behaviors, to learning behaviors from scalar feedback. While a detailed comparison of the different approaches to robot learning is beyond the scope of this paper (see [17]), it is arguable that in the short term, robots are going to be dependent on human trainers for much of their learning. S. MAHADEVAN, G. THEOCHAROUS, N. KHALEELI Specifically, a pragmatic approach to robot learning is one where a human designer provides the basic ingredients of the solution (e.g. the overall control architec- ture), with the missing components being filled in by additional training. Also, approaches involving considerable trial-and-error, such as reinforcement learning [25], are difficult to use in many circumstances, because they require long training times, or because they expose the robot to dangerous situations. For these reasons, we adopt the framework of supervised learning, where a human trainer provides the robot with labeled examples of the desired concept. Supervised concept learning from labeled examples is probably the most well-studied form of learning [20]. Among the most successful approaches are decision trees [23] and neural networks [19]. Concept learning in robotics is an extremely challenging problem, for several reasons. Sensory data is often very high-dimensional (e.g. even a coarsely subsampled image can contain millions of pixels), noisy due to specularities and other irregularities, and typically data collection requires the robot to move to different parts of its environment. Under these con- ditions, it seems clear that some form of a priori knowledge or bias is necessary for robots to be able to successfully learn interesting concepts. In this paper, we investigate two general approaches to bias sensory concept learning for mobile robots. The first is based on spatial decomposition of the sensor representation. The idea here is to partition a high-dimensional sensor represen- tation, such as a local occupancy grid or a visual image, into multiple quadrants, and learn independently from each quadrant. The second form of bias investigated here is to learn multiple concepts using a shared representation. We investigate the effectiveness of these two approaches on two realistic tasks, navigation and recy- cling. Both these tasks are studied on a real robot called PAVLOV (see Figure 1). In both problems, we use a standardized function approximator, in the form of a feedforward neural net, to represent concepts, although we believe the bias strategies studied here would be applicable to other approximators (e.g. decision trees or instance-based methods). In the navigation task, PAVLOV is required to traverse across an entire floor of the engineering building (see Figure 10). The navigational system uses a hybrid two-layered architecture, combining a probabilistic planning and execution layer with a reactive behavior-based layer. The planning layer requires the robot to map sensory values into high-level features, such as "doors" and "openings". These observations are used in state estimation to localize the robot, and are critical to successful navigation despite noisy sensing and actions. We study how PAVLOV can be trained to recognize these features from local occupancy grid data. We also show that spatial decomposition and multiple category learning provide a relatively rapid training phase. In the recycling task, PAVLOV is required to find items of trash (e.g. soda cans and other litter) and deposit them in a specified trash receptacle. The trash receptacles are color coded, to make recognition easier. Here, we study how PAVLOV can be trained to recognize and find trash receptacles from color images. The data RAPID CONCEPT LEARNING FOR MOBILE ROBOTS 3 is very high dimensional, but once again, spatial decomposition and multi-category learning are able to sufficiently constrain the hypothesis space to yield fast learning. The rest of the paper is organized as follows. We begin in Section 2 by describing the two robotics tasks where we investigated sensory concept learning. Section 3 describes the two general forms of bias, decomposition and sharing, used to make the concept learning problem tractable. Section 4 describes the experimental results obtained on a real robot platform. Section 5 discusses the limitations of our approach and proposes some directions for further work. Section 6 discusses some related work. Finally, Section 7 summarizes the paper. 2. Two Example Tasks We begin by describing the real robot testbed, followed by a discussion of two tasks involving learning sensory concepts from high-dimensional sensor data. The philosophy adopted in this work is that the human designer specifies most of the control architecture for solving the task, and the purpose of sensory concept learning is to fill in some details of the controller. 2.1. PAVLOV: A Real Robot Figure shows our robot PAVLOV 1 , a Nomad 200 mobile robot base, which was used in the experiments described below. The sensors used on PAVLOV include ultrasound sonar and infra-red (IR) sensors, arranged radially in a ring. Two sets of bumper switches are also provided. In addition, PAVLOV has a color camera and frame-grabber. Communication is provided using a wireless Ethernet system, although most of the experiments reported in this paper were run onboard the robot's Pentium processor. Figure 1. The experiments were carried out on PAVLOV, a Nomad 200 platform. 4 S. MAHADEVAN, G. THEOCHAROUS, N. KHALEELI 2.2. Navigation Robot navigation is a very well-studied topic [1]. However, it continues to be an active topic for research since there is much room for improvement in current systems. Navigation is challenging because it requires dealing with significant sensor and actuator errors (e.g. sonar is prone to numerous specular errors, and odometry is also unreliable due to wheel slippage, uneven floors, etc. We will be using a navigation system based on a probabilistic framework, formally called partially-observable Markov decision processes (POMDP's) [4], [13], [21]. This framework uses an explicit probabilistic model of actuator and sensor uncer- tainty, which allows a robot to maintain belief estimates of its location in its envi- ronment. The POMDP approach uses a state estimation procedure that takes into account both sensor and actuator uncertainty to determine the approximate location of the robot. This state estimation procedure is more powerful than traditional state estimators, such as Kalman filters [14], because it can represent discontinuous distributions, such as when the robot believes it could be in either a north-south corridor or an adjacent east-west corridor. For state estimation using POMDP's, the robot must map the current sensor values into a few high level observations. In particular, in our system, the robot generates 4 observations (one for each direction). Each observation can be one of four possibilities: door, wall, opening, or undefined. These observations are generated from a local occupancy grid representation computed by integrating over multiple sonar readings. Figure 2 illustrates the navigation system onboard PAVLOV, which combines a high level planner with a reactive layer. The route planner and execution system used is novel in that it uses a discrete-event probabilistic model, unlike previous approaches which use a discrete-time model. However, as the focus of this paper is on learning the feature detectors, we restrict the presentation here to explaining the use of feature detectors in state estimation, and refer the reader to other sources for details of the navigation system [11], [18]. The robot maintains at every step a belief state, which is a discrete probability distribution on the underlying state space (e.g. in our environment, the belief state is a 1200-dimensional vector). If the current state distribution is ff prior , the state distribution ff post , after the execution of an abstract action a, is given by 2 ff post scale This updated state distribution now serves as ff prior when the state distribution is updated to ff post , after an abstract observation ff post scale O(ojs)ff prior (s); 8s 2 S (2) In both updates, scale is a normalization constant that ensures that RAPID CONCEPT LEARNING FOR MOBILE ROBOTS 5 ff post This is necessary since not every action is defined in every state (for example, the action go-forward is not defined in states where the robot is facing a wall). 2.3. Abstract Observations In each state, the robot is able to make an abstract observation. This is facilitated through the modeling of four virtual sensors that can perceive features in the nominal directions front, left, back and right. Each sensor is capable of determining if a percept is a wall, an opening, a door or if it is undefined. An abstract observation is a combination of the percepts in each direction, and thus there are 256 possible abstract observations. The observation model specifies, for each state and action, the probability that a particular observation will be made. Denote the set of virtual sensors by I and the set of features that sensor i 2 I can report on by Q(i). The sensor model is specified by the probabilities v i (f js) for all i 2 I , f 2 Q(i), and s 2 S, encoding the sensor uncertainty. v i is the probability with which sensor i reports feature f in state s. An observation o is the aggregate of the reports from each sensor. This is not explicitly represented. We calculate only the observation probability. Thus, if sensor i reports feature f , then Y i2I Given the state, this assumes sensor reports from different sensors are independent.0000000000000000000000000000000000000000000000000000000000000000111111111111111111111111111111111111111111111111111111111111111111111111 Sensor Reports Grids Raw Sensor Values Motor Commands Action Reports Action Commands Neural Net Feature Detectors Layer Behavior-based Layer Planning Figure 2. A hybrid declarative-reactive architecture for robot navigation. The neural net feature detectors (shaded box) are trained using spatial decomposition and multi-task learning. 6 S. MAHADEVAN, G. THEOCHAROUS, N. KHALEELI 2.4. Recycling The second task we study is one where the robot has to find and pick up litter lying on the floor (e.g. soda cans and other junk) and deposit it in a colored trash receptacle (see Figure 3). This task involves several component abilities, such as locating and picking up the trash, and also subservient behaviors (such as avoiding obstacles etc. However, for the purposes of this paper, we will mainly focus on the task of detecting a trash can from the current camera image, and moving the robot till it is located adjacent to the trash can. Figure 3. Image of a trash can, which is color coded to facilitate recognition (this can is colored yellow). Avoid Motors/Turret Turn Camera Sensors Avoid Bump Figure 4. Behavior-based architecture for recycling task. The focus of sensory concept learning here is to improve "camera turn" behavior by learning how to detect and move towards trash cans. RAPID CONCEPT LEARNING FOR MOBILE ROBOTS 7 The recycling task is accomplished using a behavior-based architecture [2], as illustrated in Figure 4. Only one of the behaviors, "camera turn" is improved by the sensory concept learning methods described here, in particular, by learning how to detect and move towards the trash can. The other behaviors implement a collection of obstacle avoidance algorithms, which are not learned. 3. Accelerating sensory concept learning Learning sensory concepts is difficult because the data is often very high-dimensional and noisy. The number of instances is often also limited, since data collection requires running the robot around. In order to learn useful concepts, under these conditions, requires using some appropriate bias [20] to constrain the set of possible hypotheses 3 The study of bias is of paramount importance to machine learning, and some researchers have attempted a taxonomy of different type of bias (e.g. see [24]). Among the main categories of bias studied in machine learning are hypothesis space bias (which rules out certain hypotheses), and preference bias which ranks one hypotheses over another (e.g. prefer shallower decision trees over deeper ones). In the context of robotics, the ALVINN system [22] for autonomous driving is a good example of the judicious use of hypotheses bias to speed convergence. Here, for every human provided example, a dozen or so synthetic examples are constructed by scaling and rotating the image input to the net, for which the desired output is computed using a known pursuit steering model. We present below two ways of accelerating sensory concept learning, which can also be viewed as a type of hypotheses space biases. 3.1. Spatial Decomposition The sensory state space in both tasks described above is huge (of the order of several hundred real-valued inputs). The number of training examples available is quite limited, e.g. on the order of a few hundred at most. How is it possible to learn a complex function from such a large state space, with so little data? We use two general approaches to decompose the overall function learning problem. The first idea is simple: partition the state into several distinct regions, and learn subfunctions for each region. The idea is illustrated in Figure 5. This idea is used in the navigation domain to train four separate feature detectors, one each for the front and back quadrants of the local occupancy grid, and one each for the left and right quadrants. There are two advantages of such a decomposition: each image generates four distinct training examples, and the input size is halved from the original input (e.g. in the navigation domain, the number of inputs is 512 rather than 1024). 8 S. MAHADEVAN, G. THEOCHAROUS, N. KHALEELI 3.2. Multi-class Learning The second strategy used in our work to speed sensory concept learning is to learn multiple categories using a shared structure. This idea is fairly well-known in neural nets, where the tradeoff between using multiple single output neural nets vs. one multi-output neural net has been well studied. Work by Caruana [3] shows that even when the goal is to learn a single concept, it helps to use a multi-output net to learn related concepts. Figure 6 illustrates the basic idea. In the recycling domain, for example, the robot learns not just the concept of "trash can", but also whether the object is "near" or "far", on the "left" or on the "right". Simultaneously learning these related concepts results in better performance, as we will show below. 4. Experimental Results The experiments described below were conducted over a period of several months on our real robot PAVLOV, either inside the laboratory (for recycling) or in the corridors (for navigation). We first present the results for the navigation task, and subsequently describe the results for the recycling task. 4.1. Learning Feature Detectors for Navigation Given that the walls in our environment were fairly smooth, we found that sonars were prone to specular reflections in a majority of the environment. This made it difficult to create hard-coded feature detectors for recognizing sonar signatures. We show below that using an artificial neural network produced more accurate and consistent results. Not only was it easy to implement and train, but it is also possible to port it to other environments and add new features. Figure 7 shows the neural net used in feature detection. The net was trained using the quickprop method [8], an optimized variant of the backpropagation algorithm. local occupancy grids were collected by running the robot through the hallways. Each local occupancy grid was then used to produce 4 training patterns. The neural net was trained on 872 hand labeled examples. Since all sensors predict the same set of features, it was only necessary to learn one set of weights. Figure 8 shows the learning curve for the neural net, using batch update. Starting off with a set of random weights, the total error over all training examples converged to an acceptable range (! 1) within about training epochs. A separate set of data, with 380 labeled patterns, was used to test the net. This would be the approximately the number of examples encountered by the robot, as it navigated the loop in the Electrical Engineering department (nodes 3-4-5-6 in Figure 10). Feature prediction is accomplished by using the output with the maximum value. Out of the 380 test examples, the neural net correctly predicts features for 322, leading to an accuracy of 85%. RAPID CONCEPT LEARNING FOR MOBILE ROBOTS 900000000000000000000000000000000011111111111111111111111111111111111111111111000000000000000000000000000000000000111111111111111111111111111111111111 sensor representation original function front back left right Figure 5. Spatial decomposition of the original sensory state helps speed learning sensory concepts. Here, the original sensory space is decomposed into a pair of two disjoint quadrants.00110011001101010011 001100000000000000000000000001111111111111111111111111000000000000000000000000000000000000000000000000001111111111111111111111111111111111111111111110000011111000000000000000000000000011111111111111111111111110000011111 00000000000000000000000000000011111111111111111111111110000000000000000000000000111111111111111111111111100000000000000000000000001111111111111111111111111000000111111000000111111000000111111 Inputs Concept1 Concept2 Concept3 Concept4 Figure 6. Learning multiple concepts simultaneously using a shared representation can speed sensory learning. S. MAHADEVAN, G. THEOCHAROUS, N. KHALEELI Front Back Left Right Hidden Layer Output Layer Input Layer local occupancy grid (32x32) Door Wall Opening Undefined Figure 7. A local occupancy grid map, which is decomposed into four overlapping quadrants (left, right, top, bottom), each of which is input to a neural net feature detector. The output of the net is a multi-class label estimating the likelihood of each possible observation (door, opening, wall, or undefined). The net is trained on manually labeled real data.20060010000 Total Epochs curve for neural net feature detector Figure 8. Learning curve for training neural net to recognize features. The net is trained on 872 hand labeled examples using quickprop. RAPID CONCEPT LEARNING FOR MOBILE ROBOTS 11 Opening) (Right Opening) (Back Opening) (Back Opening) Wall) (Right Opening) (Back Opening) Wall) (Right Wall) (Back Opening) Wall) (Right Wall) (Back Opening) Wall) (Right Wall) (Back Opening) Wall) (Right Wall) (a) (c) (b) (d) Figure 9. Sample local occupancy grids generated over an actual run, with observations output by the trained neural net. Despite significant sensor noise, the net is able to produce fairly reliable observations. S. MAHADEVAN, G. THEOCHAROUS, N. KHALEELI Figure 9 illustrates the variation in observation data, generated during an actual run. In these occupancy grids, free space is represented by white, while black represents occupied space. Gray areas indicate that occupancy information is un- known. The figures are labeled with the virtual sensors and corresponding features, as predicted by the neural net. Specular reflections occur when a sonar pulse hits a smooth, flat surface angled obliquely to the transducer. The possibility exists that the sonar pulse will reflect away from the sensor, and undergo multiple reflections before it is received by the sensor. As a result, the sensor registers a range that is substantially larger than the actual range. In the occupancy grids, this results in a physically occupied region having a low occupancy probability. In Figure 9(a) where the specularities are relatively insignificant, the neural net does an accurate job of predicting the features. Effects of the specularities are noticeable in Figure 9(b) and Figure 9(c). In Figure 9(b) the neural net is able to predict a wall on the left, although it has been almost totally obscured by specular reflections. The occupancy grid in Figure 9(c) shows some bleed-through of the sonars. In both examples, the neural net correctly predicts the high level features. Figure 9(e) and Figure 9(f) are examples of occupancy grids where the effects of the specularities become very noticeable. In these examples specularities dominate, almost totally wiping out any useful information, yet the neural net is still able to correctly predict features. From the presented examples, it is apparent that the neural net can robustly predict features in a highly specular environment. Testing the neural net on an unseen set of labeled data reveals that it is able to correctly predict 85% of the features. In addition, although examples have not been presented, the neural net is able to accurately predict features even when the robot is not approximately oriented along one of the allowed compass directions. The navigation system was tested by running the robot over the entire floor of the engineering building over a period of several months (see Figure 10). The figure also shows an odometric trace of a particular navigation run, which demonstrates that despite significant odometric and sensor errors, the robot is still able to complete the task. 4.2. Learning to Find Trash Cans We now present the experimental findings from the recycling task. In order to implement a similar neural network approach, we first took various snapshots of the trash can from different angles and distances using the on-board camera of PAVLOV. The images (100x100 color images) were labeled as to the distance and orientation of the trash can. Six boolean variables were used to label the images (front, left, right, far, near, very-near). The inputs to our neural network were pre-processed selected pixels from the 100X100 images, and the outputs were the six boolean variables. The RGB values of the colored images were transformed into HSI values (Hue, Saturation, Intensity) which are better representatives of true color value because they are more invariant RAPID CONCEPT LEARNING FOR MOBILE ROBOTS 13 Electrical Engineering Department Main Office Engineering Computing Computer Science Department 3.25 mm Faculty Office Robot Lab18 Y (meters) Odometric plot of three successive runs on PAVLOV in the EE department Figure 10. The 3rd floor of the engineering building was used to test the effectiveness of the feature detectors for navigation. The bottom figure shows an odometric trace of a run on PAVLOV, showing the robot starting at node 1, doing the loop (3-4-5-6), and returning to node 1. The robot repeated this task three times, and succeeded despite significant odometric errors. 14 S. MAHADEVAN, G. THEOCHAROUS, N. KHALEELI to light variations [10]. Using an image processing program we identified the HSI values of the yellow color and based on those values we thresholded the images into black and white. We then sub-sampled the images into 400 pixels so that we could have a smaller network with far fewer inputs. The sub-sampling was done by selecting one pixel in every five. Figure 11 shows the neural net architecture chosen for the recycling task. Figure shows some sample images, with the output generated by the trained neural net. The neural net produces a six element vector as its output, with 3 bits indicating the direction of the trash can (left, front, or right), and 3 bits indicating the distance (far, near, very near). The figure shows only the output values that were close to 1. Note that the net can generate a combination of two categories (e.g. near and very-near), or even sometimes a contradictory labeling (e.g. far/near). In such cases, the camera turn behavior simply chooses one of the labels, and proceeds with capturing subsequent images, which will eventually resolve the situation (this is shown in the experiments below). Figure 13 shows the learning curve for training the trash can net. Figure 14 shows the experimental setup used to test the effectiveness of the trash can finder. A single yellow colored trash can was placed in the lab at four different locations. In each case, the robot was started at the same location, and its route measured until it stopped adjacent to the trash can (and announced that it had found the trash can). 4 Figure 15 and Figure 16 show several sample trajectories of the robot as it tried to find the trash can. In all cases, the robot eventually finds the trash can, although it takes noticeably longer when the trash can is not directly observable from the starting position. 5. Limitations of the Approach The results presented above suggest that high-dimensional sensory concepts can be learned from limited training examples, provided that a human designer carefully structures the overall learning task. This approach clearly has some definite strengths, as well as some key limitations. ffl Need for a teacher: Supervised concept learning depends on a human teacher for providing labeled examples of the desired target concept. Previous work on systems such as ALVINN [22] has clearly demonstrated that there are interesting tasks where examples can be easily collected. Similarly, for the navigation and recycling task, we have found that collecting and labeling examples to be a fairly easy (although somewhat tedious) task. Nevertheless, this approach could not be easily used in domains where it is difficult for a human teacher to find a sufficiently diverse collection of examples. ffl Filling in details of a pre-specified architecture: The approach taken in this paper assumes that the designer has already pre-specified much of the overall control structure for solving the problem. The purpose of learning is to complete RAPID CONCEPT LEARNING FOR MOBILE ROBOTS 15 inputs 6 outputs hidden units RIGHT FRONT FAR VERY NEAR Figure 11. A neural net trained to detect trash cans. a b c Figure 12. Sample images with the output labels generated by the neural net a: front, near. b: front, far. c: left, front, far. d: right, far. e: left, near, very-near. f: front, very-near. S. MAHADEVAN, G. THEOCHAROUS, N. KHALEELI20601001401800 20 40 Epochs Comparing Multi-Output Learning vs. Single Output Learning 'multi-output.err' 'very-near.err' Figure 13. This graph compares the training time for a multi-output net vs. training a set of single output nets. Although the multi-output net is slower to converge, it performed better on the test data. a few missing pieces of this solution. In the navigation task, for example, the feature detectors are all that is learned, since the overall planner, reactive behaviors, and state estimator are pre-programmed. Obviously, this places a somewhat large burden on the human designer. ffl Decomposable functions: The sensory concepts being learned in the two tasks were decomposable in some interesting way (either the input or the output space could be partitioned). We believe many interesting concepts that robots need to learn have spatial regularity of some sort that can be exploited to facilitate learning. 6. Related Work This research builds on a distinguished history of prior work on concept learning from examples, both in machine learning [20] and in robot learning [5], [9]. Here, we focus primarily on the latter work, and contrast some recent neural-net based approaches with decision-tree based studies. ALVINN [22] uses a feedforward neural net to learn a steering behavior from labeled training examples, collected from actual human drivers. As noted earlier, ALVINN exploits a pursuit model of steering to synthesize new examples to speed RAPID CONCEPT LEARNING FOR MOBILE ROBOTS 17 Figure 14. Environmental setup for finding trash cans. learning. ALVINN differs from our work in that it directly learns a policy, whereas in our case the robot learns only feature detectors and recognizers. We believe that directly learning an entire policy is quite difficult, in general. In fact, in a subsequent across-the-country experiment, the direct policy learning approach was rejected in favor of a simpler feature-based approach similar to our work (except the templates were 1-dimensional rather than 2-dimensional, as in our work). Thrun and Mitchell [27] propose a lifelong learning approach, which extends the supervised neural-net learning framework to handle transfer across related tasks. Their approach is based on finding invariances across related functions. For ex- ample, given the task of recognizing many objects using the same camera, invariances based on scaling, rotation, and image intensity can be exploited to speed up learning. Their work is complementary to ours, in that we are focusing on rapid within-task learning, and the invariants approach could be easily combined with the partioning and multi-class approach described here. Such studies can be contrasted with those using decision trees. For example, Tan [26] developed an ID3 decision-tree based algorithm for learning strategies for picking up objects, based on perceived geometric attributes of the object, such as its height and shape. Salganicoff et al. [15] extended the decision-tree approach for learning grasping to an active learning context, where the robotic system could itself acquire new examples through exploration. In general, the decision tree approaches seem more applicable when the data is not high-dimensional (in both the system just cited, the number of input dimensions is S. MAHADEVAN, G. THEOCHAROUS, N. KHALEELI -1.4 -1.2 -0.4 -0.20.2 South East Navigation to lab position A "pose.a1" "pose.a2" "pose.a3" -0.3 -0.2 South East Navigation to lab position B "pose.b1" "pose.b2" "pose.b3" Figure 15. Three successful traces (starting at 0,0) of the robot navigating to the trash can, placed in positions A and B. In both positions A and B, the trash can was directly observable from the robot starting position. South East Navigation to lab position C "pose.c1" "pose.c2" "pose.c3" South East Navigation to lab position D "pose.d1" Figure 16. Results for learning with trash can in position C and D. The top figure shows three successful traces (starting at 0,0) of the robot navigating to lab position C. Note that in pose.c2 trace the robot temporarily loses the trash can but eventually gets back on track. The bottom figure shows a successful run when the trash can is in position D, which is initially unobservable to the robot. S. MAHADEVAN, G. THEOCHAROUS, N. KHALEELI generally less than 10). By contrast, in our work as well as in the ALVINN system, the input data has several hundred real-valued input variables, making it difficult to employ a top decision-tree type approach. The advantage, however, of using decision trees is that the learned knowledge can be easily converted into symbolic rules, a process that is much more difficult to do in the case of a neural net. Symbolic learning methods have also been investigated for sensory concept learn- ing. Klingspor et al. [12] describe a relational learning algorithm called GRDT, which infers a symbolic concept description (e.g. the concept thru door) by generalizing user labeled training instances of a sequence of sensor values. A hypothesis space bias is specified by the user in the form of a grammar, which restrict possible generalizations. A strength of the GRDT algorithm is that it can learn hierarchical concept descriptions. However, a weakness of this approach is that it relies on using a logical description of the overall control strategy (as opposed to using a procedural reactive/declarative structure). Logical representations incur a computational cost in actual use, and their effectiveness in actual real-time robotics applications has not been encouraging. 7. Summary This paper investigates how mobile robots can acquire useful sensory concepts from high-dimensional and noisy training data. The paper proposes two strategies for speeding up learning, based on decomposing the sensory input space, and learning multiple concepts simultaneously using a shared representation. The effectiveness of these strategies was studied in two tasks: learning feature detectors for probabilistic navigation and learning to recognize visual objects for recycling. A detailed experimental study was carried out using a Nomad 200 real robot testbed called PAVLOV. The results suggest that the strategies provide sufficient bias to make it feasible to learn high-dimensional concepts from limited training data. Acknowledgements This research was supported in part by an NSF CAREER award grant No. IRI- 9501852. The authors wish to acknowledge the support of the University of South Florida Computer Science and Engineering department, where much of this research was conducted. We thank Lynn Ryan for her detailed comments on a draft of this paper. Notes 1. PAVLOV is an acronym for Programmable Autonomous Vehicle for Learning Optimal Values. 2. In actuality, the state estimation procedure is more complex since we use an event-based semi-Markov model to represent temporally extended actions. However, for the purposes of this paper, we are simplifying the presentation. RAPID CONCEPT LEARNING FOR MOBILE ROBOTS 21 3. Bias is generally defined as any criterion for selecting one generalization over another, other than strict consistency with the training set. It is easy to show that bias-free learning is impossible, and would amount to rote learning. 4. Although we do not discuss the details here, the robot employs a further processing phase to extract the rough geometrical aligment of the trash can opening in order to drop items inside it. --R Navigating Mobile Robots. A robust layered control system for a mobile robot. Multitask learning: A knowledge-based source of inductive bias Acting under uncertainty: Discrete bayesian models for mobile-robot navigation Robot Learning. Introduction to the special issue on learning autonomous robots. Robotics in Service. Robot Learning. Fundamentals of Digital Image Processing. A robust robot navigation architecture using partially observable semi-markov decision processes Learning concepts from sensor data. A robot navigation architecture based on partially observable markov decision process models. Fast vision-guided mobile robot navigation using model-based reasoning and prediction of uncertainties Machine learning for robots: A comparison of different paradigms. Mobile robot navigation using discrete-event markov decision process models Machine Learning. An office-navigating robot Neural network based autonomous navigation. Induction of decision trees. Inductive learning from preclassified training examples. Reinforcement Learning: An Introduction. Learning one more thing. --TR --CTR B. L. Boada , D. Blanco , L. Moreno, Symbolic Place Recognition in Voronoi-Based Maps by Using Hidden Markov Models, Journal of Intelligent and Robotic Systems, v.39 n.2, p.173-197, February 2004

robot learning;neural networks;concept learning

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Power balance and apportionment algorithms for the United States Congress.

We measure the performance, in the task of apportioning the Congress of the United States, of an algorithm combining a heuristic-driven (simulated annealing) search with an exact-computation dynamic programming evaluation of the apportionments visited in the search. We compare this with the actual algorithm currently used in the United States to apportion Congress, and with a number of other algorithms that have been proposed. We conclude that on every set of census data in this country's history, the heuristic-driven apportionment provably yields far fairer apportionments than those of any of the other algorithm considered, including the algorithm currently used by the United States for Congressional apportionment.

1. MOTIVATION AND OVERVIEW How should the seats in the House of Representatives of the United States be allocated among the states? The Constitution stipulates only that "Representa- tives shall be apportioned among the several states according to their respective numbers, counting the whole numbers of persons in each :." The obvious implementation of this requirement would almost always yield fractional numbers of seats. The issue of how to achieve fair integer seat allocations has been controversial in this country virtually since its founding. In fact, many of the apportionment algorithms we will discuss have been proposed and debated by famous historical figures, including John Quincy Adams, Alexander Hamilton, Thomas Jefferson, and Daniel Webster. The debate is far from over. In fact, the relative fairness of two of the algorithms we will discuss in this paper was argued before the Supreme Court in 1992 [Supreme 1992]. 1 We propose an apportionment method consisting of a simulated-annealing search that is aimed at maximizing the fairness of the resultant apportionments. Even though the complexity of the algorithm is high, our implementation shows that the method is feasible for the cases of interest. In- deed, we have been able to run it on the data conducted during all the census years in US history and the results are conclusive: In all cases our method was provably superior with respect to widely agreed fairness criteria to the most prominent apportionment algorithms that have been used or proposed earlier. Balinski and Young [1982] (see also [Balinski and Young 1985]) performed a detailed comparative study, for six historical algorithms of Congressional apportion- ment, of the degree to which the algorithms' allocations matched states' "quotas," i.e., their portion of the population times the House size. Mann and Shapley [Mann and Shapley 1960; Mann and Shapley 1962] and others studied, for the actual used- in-Congress seat allocations, the power indices (in the Electoral College) of each state. This paper attempts to combine the strengths of these two research lines. In particular, we agree with Balinski and Young both that allocations should be "fair," and that, in light of 200 years of debate (colorfully recounted and analyzed by Balinski and Young [1982]), obtaining new insights into the merits and weaknesses of the six historical algorithms should be a priority. On the other hand, many feel that "fairness" should be defined by a tight match between power and quotas, rather than between allocations and quotas. Our feeling is very much in harmony with modern political science theory, where it is widely recognized that allocations Briefly put, the Supreme Court ruled that Congress acted within its authority in choosing the currently used algorithm (the "Huntington-Hill Method"). However, the Supreme Court's decision left open the possibility that Congress would be acting equally within its authority if it chose to adopt some other algorithm. The court ruled that "the constitutional framework.delegate[s] to Congress a measure of discretion broader than that accorded to the States [in terms of choosing how to apportion], and Congress's apparently good-faith decision to adopt the [Huntington-]Hill Method commands far more deference, particularly as it was made after decades of experience, experimentation, and debate, was supported by independent scholars, and has been accepted for half a century [Supreme 1992]." Regarding the decision, we mention only that the power balancing issues discussed in this paper were not brought before the court, but that nonetheless the experimental results of this paper suggest that among the two algorithms being discussed by the court in the case, the Huntington-Hill algorithm in fact gives fairer apportionments in terms of power balancing. Power Balance and Apportionment Algorithms \Delta 3 do not necessarily directly correspond to power, and that power indices (see the detailed definitions later in this paper) provide a potentially more accurate gauge of power (see, e.g., the discussions in [Shapley 1981; Riker and Ordeshook 1973]). 2 So, in this light, we repeat the comparative study of Balinski and Young, but we replace allocation comparisons with power index comparisons. Of course, the historical algorithms were designed (in part-the full story is more complex and political, and indeed led to the first presidential veto (see [Balinski and Young 1982])) to achieve some degree of harmony and fairness between allocations and quotas. This is not surprising, given that power indices had not yet been invented. However, given that in this study our comparisons are based on power indices, it seems natural to add to the six historical algorithms 3 an algorithm tailored to achieve harmony between power indices and quotas. We have used a heuristic based on the simulated annealing paradigm (see [Aarts and Korst 1989; Metropolis et al. 1953; Kirkpatrick et al. 1983]), which finds an apportionment by seeking to achieve a local minimum of the distance between normalized power indices and quotas (the attribute "local" is with respect to a natural neighborhood relation between apportionments). This heuristic yields results that are fairer to those obtained by all the historical algorithms. We report in Section 3 the results for the last census, 1990, but we have obtained similar results for all the census years, 1790, 1800, . , 1980, and 1990. The function class #P (first defined by Valiant [Valiant 1979a; Valiant 1979b]) is the counting version of NP. #P is the class of all functions f such that, for some nondeterministic polynomial-time Turing machine N , for all inputs x it holds that f(x) is the number of accepting computation paths of N(x). One problem in studying power indices is that power indices are typically #P-complete [Prasad and Kelly 1990; Garey and Johnson 1979] and, consequently, we perform a combinatorial search that invokes at each iteration numerous #P-complete computations. Fortunately, a dynamic programming approach, first proposed by Mann and Shapley [1962], yields a pseudo-polynomial algorithm for computing the power indices, i.e., an algorithm whose running time is polynomial in the size of the House and in the number of states. Since these quantities have had reasonable values throughout US history (the maximum values have been 435 and 50, which are also the current values), we have been able to exactly compute the needed power indices. 2 For those unfamiliar with why allocations may not correspond to power, consider the following typical motivating example. Suppose we have states A, B, and C with 6, 2, and 2 votes respectively. Note that though between them states B and C have 40 per cent of the seat allocation, nonetheless it is the case that in a majority-rule vote on some polarizing issue on which states have differing interests and so their delegations vote as blocs, B and C have no power at all as A by itself is a majority. 3 By the use of "the" in the context "the six historical algorithms" we do not mean to suggest that no other algorithms have been proposed. Other algorithms have indeed been proposed (e.g., the algorithm Condorcet suggested in 1793 ([Condorcet 1847], see [Balinski and Young 1982, p. 63])). However, these six algorithms have been the key contenders in the apportionment discussion in the United States. 2. DISCUSSION In this section, we justify a number of decisions made in designing this study, and we describe in more detail the background of the study. 2.1 Study Design Our computer program takes as its input a list of states, hS tions, h, and some other parameters regarding random number generation and the implementation of the simulated annealing algorithm. The program computes, for each state S i , the appropriate quota We have considered the Banzhaf power index and the Shapley-Shubik power index. They represent the most widely used quantitative ways to measure the power of a player (or a state, in our parlance) in a voting game. It is well-known that the two indices can, at least on artificially constructed examples, differ sharply (see, e.g., [Straffin 1978] and the references therein, and [Shapley 1981; Straffin 1977]). We do not in this paper seek to resolve the broader question of whether power indices are "correct" or "appealing" measures of power, or whether they thus should be used to shape voting/apportionment systems (those are more issues of political science than of experimental algorithms). Though there is much discussion in the literature on the naturalness and usefulness of power indices, our study simply seeks to provide-to those who do place value in power indices-quantitative information on the comparative power distributions for the historical apportionment algorithms and for simulated annealing. We hope that this information will also be useful in the political science debate on the more general issue of choosing the "right" apportionment scheme, and in Section 4 we mention the directions this information suggests. The Banzhaf power index (see, e.g., the detailed discussion in [Dubey and Shapley 1979]) is defined in terms of winning coalitions. A winning coalition is a subset S of states such that the sum of the votes of the states in S is larger than half of the total number of votes (which, of course, is equal to the House size). A state i is critical for a coalition S if (a) i 2 S, (b) S is a winning coalition, and (c) S \Gamma fig is not a winning coalition. The Banzhaf power index of a state i is, by definition, the probability that for a randomly chosen coalition S (that is, for a coalition chosen randomly from the set of all possible coalitions, not just from the set of winning coalitions), it holds that i is critical for S. The Shapley-Shubik power index [Shapley and Shubik 1954] is defined in terms of linear orderings of states, which, intuitively, represent the extent to which the states are interested in passing some bill (the first state in the ordering is the strongest supporter of the bill, while the last one is the strongest opponent). Let be such an ordering. A state i is a pivot for the ordering - if it appears in some position k and i is critical for the coalition formed by the first k states (in other words, the pivot of an ordering is the state whose joining turns the developing coalition into a winning coalition). The Shapley-Shubik power index of a state i is the probability that for a randomly chosen ordering - it holds that i is the pivot of -. Power Balance and Apportionment Algorithms \Delta 5 In order to calculate the closeness of the power indices to the quota, we have to normalize them. That is, if pow i is the power index (Banzhaf or Shapley-Shubik) of state i, we define normpow where, recall, h is the House size. This is the normalization used by Mann and Shapley [1960]. For the rest of this discussion, by power of a state we will mean its normalized power index (Banzhaf or Shapley-Shubik). Note that h, and thus these normalized powers make it easier to compare power-to-quota closeness. 4 The main metric that we have used to evaluate the power-to-quota closeness is the L 2 metrics applied on proportions as defined below. L 2 on proportions: 0-i!n Although other measures can be used, we feel the L 2 on proportions to be the most relevant one. We have performed experiments also with respect to the L 1 metric on differences and proportions and L 2 on differences, and in all cases we reached the same conclusions. For completeness, here are the definitions of these three other metrics. L 1 on differences: 0-i!n L 2 on differences: 0-i!n L 1 on proportions: 0-i!n The contrast between the "difference" and the "proportion" measures can be dramatic. For example, suppose states A and B have quotas 1.5 and 40.5 but are respectively given 2 and 41 votes. The "differences" measures treat the two states similarly, but in the "proportions" measures, state A is viewed as being far more out of balance than state B. However, we repeat, the conclusions of our investigation remain the same, irrespective of the metric used. The program apportions the votes via the six historical algorithms-Adams, Dean, Huntington-Hill (we abbreviate this "H-Hill" in our tables; the method is also sometimes referred to in the literature simply as the Hill Method), Webster (which is sometimes also referred to in the literature as the Webster-Willcox Method), Jef- ferson, and Hamilton-and simulated annealing. The algorithms work as follows. Let us first discuss the algorithms of Adams, Dean, Huntington-Hill, Webster, and Jefferson. Informally, they work as follows. One partitions the non-negative real 4 We note that many researchers prefer to use the Banzhaf index without normalization. 6 \Delta L.A. Hemaspaandra et al. line into adjacent intervals, each corresponding to a number of votes. Then one seeks an integer such that if one divides each state's population by that integer and give each state the number of votes corresponding to the interval in which that value falls, the sum of votes handed out equals the desired house size. More formally, let I be such that I We say an apportionment hv 0 ,v 1 is the 5 sliding divisor apportionment with respect to I there exists a real number d such that The five historical sliding divisor algorithms are defined as follows (see also [Balinski and Young 1982]). Adams:. I I Dean:. I I Huntington-Hill:. I I Webster:. I I Jefferson:. I Note, in particular, that the Dean, Huntington-Hill, and Webster algorithms describe the intervals between, respectively, the harmonic, geometric, and arithmetic means of successive integers. The algorithm currently in use in the United States is the method of geometric means, i.e., the Huntington-Hill algorithm. An important caveat is that the Constitution requires that each state be given at least one representative. However, this artificial requirement would taint the comparative analysis of the algorithms. In this study, we do allow algorithms to assign zero votes to a state. For example, for the 1990 census data the Jefferson algorithm allocates 0 votes to Wyoming. To learn whether our ignoring the Constitution's "one vote minimum" rule affected our results, we also ran our program with the added proviso that the sliding divisor algorithms were not allowed to assign 0 votes to any state (we achieved this by setting I 1 := I 0 [ I 1 and I 0 := ;) and we obtained the same result: The simulated annealing algorithm was still the best and the Jefferson algorithm was still the worst. 5 It is not hard to see that if such exists, it is unique. Such solutions do not always exists. For example, no sliding divisor algorithm can apportion two equal states with any odd House size. However, for the five historical sliding divisor algorithms and large populations that are random and independent in their low-order bits (informally, actual populations), nonexistence of solutions is unlikely to occur, and, for historical census data, has never occurred. Power Balance and Apportionment Algorithms \Delta 7 The Hamilton algorithm is quite different from the sliding divisor algorithms described above. The Hamilton algorithm initially assigns to each state, S i , exactly votes. Then it assigns one extra vote to each of the h \Gamma states having the largest values of q that ties in the fractional part can cause this algorithm to fail to be well-defined, but, again, in practice this is unlikely to occur, and indeed never has under actual historical census data.) Finally, the simulated annealing algorithm that we propose works as follows. Let us say that two apportionments are neighbors of each other if b is obtained from a by shifting one vote from some state i to some state j, i.e., there exist i and j, i 6= j, such that apportionment (summing to h) is made in a relatively arbitrary way. (The algorithm used in the program intentionally makes an initial apportionment that gives a terrible match between powers and quotas. Thus, it is up to the rest of the simulated annealing algorithm to try to fix this poor initial assignment.) Then, repeatedly, a neighbor b of the current apportionment a is randomly generated. If b is at least as good as a (in the sense that b has a power-to-quota distance no greater than that of a), then b becomes the current apportionment. Even if the new apportionment b is worse, it still has a chance to become the current apportionment. Namely, let \Delta be the difference between the power-to-quota distance for b and the power-to-quota distance for a. Since we are in the case in which b is worse than a, \Delta is positive. Throughout this process, we consider a current temperature that is gradually decreasing (each time a given number of iterations have occurred, the current temperature is decreased by multiplying it by a certain constant that is less than one, which is called the cooling Let t be the value of the current temperature. We randomly generate a real number r. If r ! e \Gamma\Delta=t , then b is taken as the current apportionment in place of a. Observe that as the temperature t approaches the value zero, the probability that r is less than e \Gamma\Delta=t gets smaller and smaller, and thus after an initial period in which jumps between various descending hills leading to local optima are likely, the process is stabilized ("it freezes," i.e., the probability of choosing an apportionment that is worse than the current one becomes close to zero), with high probability at some local optimum (and the hope of the simulated annealing approach is that that local optimum is not too much worse in quality than the global optimum). The exact stopping criterion is explained in Section 3. The results of the simulated annealing algorithm are, of course, sensitive to the values of the initial temperature and of the cooling factor, as well as to the choice of the initial apportionment. Our intention has been not to fine-tune these values but rather to prove that the historical methods can be easily, soundly beaten by a modern computer science-based heuristic approach, on inputs drawn from real census data. As we want to emphasize the realization of this goal, we report in Section 3 the results for only one (rather arbitrary) setting of these parameters. We plan to pursue in the future a more thorough investigation of the impact of the parameter settings for the apportionment problem. We digress here to mention that this algorithm can potentially beat the greedy algorithm, 6 in which at each step we choose a better neighbor of the current ap- 6 However, for the real census data that we have used, the greedy algorithm yielded results very 8 \Delta L.A. Hemaspaandra et al. portionment until we are stuck at a local optimum. Indeed, we found a four-state instance in which the greedy algorithm gets stuck at a local optimum that is not a global optimum. In this instance, 100 representatives must be apportioned between 4 states having populations 823, 801, 105, 101. It can be seen that a = (50; 34; 14; 2) are two local optima with a being better than b (we are referring here to the Shapley-Shubik power index, which is defined below). However, starting from b and going through at-least-as-good-as b neighbors there is no way to reach a (i.e., (b; a) is not in the transitive closure of the binary relation "neighbor at-least-as-good-as b") because at some moment we have to pass through an apportionment of the form (44; ; ; ). But, it can be verified that all apportionments of this form are worse than b and thus are not reachable from b with the greedy algo- rithm. In a certain precise sense, this is the smallest such example. In particular, it is simple to prove that there are no examples such as the one above with 2 or 3 states, that is, in all 2 or 3 state cases, all locally optimal apportionments are also globally optimal. Prasad and Kelly [1990] have shown that the Banzhaf power index is #P-complete and Garey and Johnson [1979] have shown that the Shapley-Shubik power index is #P-complete, where the complexity class #P as usual denotes the counting version of NP [Valiant 1979a; Valiant 1979b]-#P is the class of all functions f such that, for some nondeterministic polynomial-time Turing machine N , for each x it holds that f(x) is the number of accepting computation paths N has on input x. Much of the past 20 years of research in theoretical computer science has been devoted to proving that NP-complete problems (and thus #P-complete functions) cannot be feasibly computed unless a wide range of implausible consequences occur (see the survey [Sipser 1992]). For example, Toda [1991] (see also [Beigel et al. 1991; Toda and Ogiwara 1992; Gupta 1995; Regan and Royer 1995]) has shown that Turing access to #P subsumes the entire polynomial hierarchy. However, a dynamic programming approach will allow us to perform an exact computation of both the Banzhaf power index and the Shapley-Shubik power index even for the relatively large inputs of the censuses involving 50 states. The dynamic programming approach for this problem was proposed by Mann and Shapley [1962]. We use the notation from the previous section. The algorithm involves, for each state constructing a matrix i C of order House size, and number of states: Also, let Q be the minimum number of votes needed for a winning coalition (i.e., 1). The entries in this matrix have the following represents the number of coalitions not containing state i and having k states whose votes sum to j. Let i C ' j;k be the number of coalitions containing k states from close to those given by the simulated annealing algorithm, in particular also beating the historical algorithms. This is true even though our implementation of the greedy algorithm accepted the first local improvement found rather than enumerating all neighbors and accepting one yielding a steepest descent. Power Balance and Apportionment Algorithms \Delta 9 whose votes sum up to k. Observe that, for Also, the following recurrence holds: (note that in the above recurrence the predecessor of j;k . To find the Banzhaf power index of state S i , we have to compute the number of coalitions for which S i is critical (and then divide it by the total number of coalitions). This value is given by These are all coalitions not containing S i that are losing but with the addition of To find the Shapley-Shubik power index of state S i , we have to compute the number of linear orderings for which S i is a pivot (and then divide it by the total number of linear orderings). This value is given by Thus, the computation of both power indices reduces to the computation of the entries of the matrix i C, which can be determined by calculating iteratively the matrices Some shortcuts can be obtained using the following observations. Note that in Equations (2.1) and (2.2), we need only the entries with 1. By looking at complementary coalitions, one can see that i C . Thus, the entries in the matrix with column index k ? (n \Gamma 1)=2 can be computed, using the above formula, from the entries with k - (n \Gamma 1)=2 and when this is done the largest row index j that we need is still Q\Gamma1. Consequently, we have to compute only a quarter of the entries in matrix i C, namely those entries with row index verifying verifying After we have computed the matrix i C for state S i , we don't have to start from scratch when passing to state S i+1 to compute i+1 C . Indeed, it is not difficult to see that i+1 C by convention the entries with negative row or column index are 0. This algorithm is polynomial in n and Q (but, of course, it is not polynomial in the length of a natural encoding of the input instance, because the reasonable encodings of Q bits). Nevertheless, since the censuses so far h - 435, we have been able to compute the exact values of the Banzhaf and Shapley-Shubik for the census years 1790, 1800, 1810, . , 1980, and State Pop. Quota SA-B SA-SS Adams Dean H-Hill Webst Jeffer Hamilt OH 10847115 19.0206 19 19 MI 9295297 16.2995 IN MO 5117073 8.9729 9 9 9 9 9 9 9 9 WI 4891769 8.5778 9 9 9 9 9 9 9 9 CO IA MS Totals 248072974 435.0000 435 435 435 435 435 435 435 435 Table 1. States, Populations, Quotas, and Apportionments: L2-of-Proportions Evaluation Power Balance and Apportionment Algorithms \Delta 11 State Quota SA-B Adams Dean H-Hill Webst Jeffer Hamilt IL 20.0437 19.8279 18.7446 19.7091 19.7097 19.7102 20.6716 19.7102 MI 16.2995 15.7744 15.7255 15.6914 15.6919 15.6924 16.6546 15.6924 GA 11.3597 11.7810 10.7599 10.7400 10.7404 10.7408 10.7208 10.7408 MA 10.5499 10.7900 9.7745 9.7569 10.7404 10.7408 10.7208 10.7408 IN 9.7218 9.8015 9.7745 9.7569 9.7573 9.7577 9.7399 9.7577 MO 8.9729 8.8152 8.7913 8.7758 8.7761 8.7765 8.7609 8.7765 WI 8.5778 8.8152 8.7913 8.7758 8.7761 8.7765 8.7609 8.7765 TN 8.5522 8.8152 8.7913 8.7758 8.7761 8.7765 7.7834 8.7765 MD 8.3844 8.8152 7.8098 7.7964 7.7967 7.7970 7.7834 7.7971 MN 7.6718 7.8308 7.8098 7.7964 7.7967 7.7970 7.7834 7.7971 LA 7.3998 7.8308 7.8098 6.8185 6.8188 6.8191 6.8074 6.8191 AL 7.0852 6.8482 6.8301 6.8185 6.8188 6.8191 6.8074 6.8191 KY 6.4622 6.8482 6.8301 5.8420 5.8422 5.8425 5.8326 5.8425 AZ 6.4270 6.8482 6.8301 5.8420 5.8422 5.8425 5.8326 5.8425 SC 6.1140 5.8671 5.8517 5.8420 5.8422 5.8425 5.8326 5.8425 CO 5.7768 5.8671 5.8517 5.8420 5.8422 5.8425 5.8326 5.8425 CT 5.7640 5.8671 5.8517 5.8420 5.8422 5.8425 5.8326 5.8425 IA 4.8691 4.8873 4.8746 4.8666 4.8668 4.8670 4.8589 4.8670 MS 4.5122 4.8873 4.8746 4.8666 4.8668 3.8925 3.8861 3.8925 KS 4.2919 3.9085 4.8746 3.8922 3.8923 3.8925 3.8861 3.8925 AR 4.1220 3.9085 3.8984 3.8922 3.8923 3.8925 3.8861 3.8925 WV 3.1449 2.9307 2.9231 2.9185 2.9186 2.9187 2.9139 2.9187 UT 3.0210 2.9307 2.9231 2.9185 2.9186 2.9187 2.9139 2.9187 NE 2.7677 2.9307 2.9231 2.9185 2.9186 2.9187 1.9423 2.9187 NM 2.6567 2.9307 2.9231 2.9185 2.9186 2.9187 1.9423 2.9187 HI 1.9433 1.9534 1.9484 1.9453 1.9454 1.9455 1.9423 1.9455 ID 1.7654 1.9534 1.9484 1.9453 1.9454 1.9455 0.9711 1.9455 RI 1.7596 1.9534 1.9484 1.9453 1.9454 1.9455 0.9711 1.9455 MT 1.4012 0.9766 1.9484 1.9453 0.9726 0.9727 0.9711 0.9727 ND 1.1201 0.9766 1.9484 0.9726 0.9726 0.9727 0.9711 0.9727 WY 0.7954 0.9766 0.9741 0.9726 0.9726 0.9727 0.0000 0.9727 Totals 435.0000 435.0000 435.0000 435.0000 435.0000 435.0000 435.0000 435.0000 Table 2. States, Quotas, and Normalized Powers: Banzhaf Power Index and L2-of-Proportions State Quota SA-SS Adams Dean H-Hill Webst Jeffer Hamilt IL 20.0437 20.0487 18.9942 19.9876 19.9858 19.9834 20.9756 19.9779 OH 19.0206 19.0005 17.9516 MI 16.2995 15.8872 15.8820 15.8399 15.8385 15.8366 16.8206 15.8324 GA 11.3597 11.8058 10.7941 10.7662 10.7653 10.7640 10.7362 10.7612 MA 10.5499 10.7974 9.7907 9.7656 10.7653 10.7640 10.7362 10.7612 IN 9.7218 9.7938 9.7907 9.7656 9.7647 9.7636 9.7384 9.7611 MO 8.9729 8.7947 8.7920 8.7695 8.7687 8.7677 8.7452 8.7654 WI 8.5778 8.7947 8.7920 8.7695 8.7687 8.7677 8.7452 8.7654 TN 8.5522 8.7947 8.7920 8.7695 8.7687 8.7677 7.7565 8.7654 MD 8.3844 8.7947 7.7977 7.7779 7.7772 7.7763 7.7565 7.7743 MN 7.6718 7.8001 7.7977 7.7779 7.7772 7.7763 7.7565 7.7743 LA 7.3998 7.8001 7.7977 6.7907 6.7902 6.7894 6.7721 6.7876 AL 7.0852 6.8100 6.8080 6.7907 6.7902 6.7894 6.7721 6.7876 KY 6.4622 6.8100 6.8080 5.8079 5.8075 5.8068 5.7921 5.8053 AZ 6.4270 6.8100 6.8080 5.8079 5.8075 5.8068 5.7921 5.8053 SC 6.1140 5.8244 5.8226 5.8079 5.8075 5.8068 5.7921 5.8053 CO 5.7768 5.8244 5.8226 5.8079 5.8075 5.8068 5.7921 5.8053 CT 5.7640 5.8244 5.8226 5.8079 5.8075 5.8068 5.7921 5.8053 OK 5.5158 5.8244 5.8226 5.8079 5.8075 4.8285 4.8164 4.8273 IA 4.8691 4.8430 4.8416 4.8295 4.8291 4.8285 4.8164 4.8273 MS 4.5122 4.8430 4.8416 4.8295 4.8291 3.8545 3.8448 3.8535 KS 4.2919 3.8660 4.8416 3.8553 3.8549 3.8545 3.8448 3.8535 AR 4.1220 3.8660 3.8649 3.8553 3.8549 3.8545 3.8448 3.8535 WV 3.1449 2.8933 2.8924 2.8853 2.8850 2.8847 2.8775 2.8840 UT 3.0210 2.8933 2.8924 2.8853 2.8850 2.8847 2.8775 2.8840 NE 2.7677 2.8933 2.8924 2.8853 2.8850 2.8847 1.9143 2.8840 NM 2.6567 2.8933 2.8924 2.8853 2.8850 2.8847 1.9143 2.8840 NV 2.1074 1.9247 1.9242 1.9194 1.9193 1.9190 1.9143 1.9186 HI 1.9433 1.9247 1.9242 1.9194 1.9193 1.9190 1.9143 1.9186 ID 1.7654 1.9247 1.9242 1.9194 1.9193 1.9190 0.9551 1.9186 RI 1.7596 1.9247 1.9242 1.9194 1.9193 1.9190 0.9551 1.9186 ND 1.1201 0.9603 1.9242 0.9577 0.9576 0.9575 0.9551 0.9573 WY 0.7954 0.9603 0.9600 0.9577 0.9576 0.9575 0.0000 0.9573 Totals 435.0000 435.0000 435.0000 435.0000 435.0000 435.0000 435.0000 435.0000 Table 3. States, Quotas, and Normalized Powers: Shapley-Shubik Power Index and L2-of- Power Balance and Apportionment Algorithms \Delta 13 Algorithm Quota to Rep-Normed Power Distance under the L2-of-Proportions Metric H-Hill 0.375372 Webst 0.387777 Hamilt 0.394220 Dean 0.438127 Adams 1.821126 Jeffer 1.906251 Table 4. Errors Between Quotas and Rep-Normed Power: Banzhaf Power Index and L2-of- Proportions Evaluation Metric. Algorithm Quota to Rep-Normed Power Distance under the L2-of-Proportions Metric H-Hill 0.382663 Webst 0.401829 Hamilt 0.411409 Dean 0.424247 Adams 1.687721 Jeffer 1.974655 Table 5. Errors Between Quotas and Rep-Normed Power: Shapley-Shubik Power Index and L2-of-Proportions Evaluation Metric. 3. RESULTS We present here the results of our program for the 1990 census. Table 1 shows the populations and quotas for the 1990 census, as well as the apportionments that result from the six historical algorithms and the simulated annealing algo- rithm, under both the Banzhaf (with column label SA-B) and the Shapley-Shubik (with column label SA-SS) power indices. Of course, the columns not reporting on the results of the simulated annealing algorithm do not need to be labeled with Banzhaf or Shapley-Shubik, as only in the simulated annealing columns is the apportionment dependent on the power index being used. Tables 2 and 3 show the normalized power indices. Tables 4 and 5 show the distance between powers and quotas according to the L2-on-proportions metric. The tables represent runs with the following parameters used in the simulated annealing algorithm: the cooling factor was 0.9, 7 the initial temperature was 100, the number of iterations per- 7 Ideally, simulated annealing should be done with a cooling factor as close to 1 as possible (as, conceptually, gentle cooling is more likely to achieve a good result). Due to the limited computing resources we had available-SUN SPARCstation 10's serving as shared cycle servers-we had to use a cooling factor that is more severe than we would have ideally chosen. Note that this computational limitation if anything degrades the quality of our simulated annealing apportionment. Yet, even with this handicap, our simulated annealing algorithm outperformed all the historical algorithms. Also, we note that if the experimental algorithms paradigm is to be successful, it is important that it be useful not just to the few people with access to supercomputers, but also to people using more modest computing facilities. We hope that this study, in which a modest workstation performed the computations showing that simulated annealing outperforms the currently used algorithm, is an example of this. 14 \Delta L.A. Hemaspaandra et al. formed at each temperature was 1000, and the stopping factor was 5. The stopping factor is used to control termination as follows. Recall that at each temperature the algorithm generates and considers CoolingIter potential swaps, and then it decreases the current temperature via Temperature := Temperature CoolingFactor. However, the algorithm also keeps track of how many potential swaps have been considered since the last swap that was performed (recall that a swap is performed either because it improves the quality of the current state, or because it degrades the quality of the current state but was randomly chosen via the temperature- based probabilistic condition discussed in Section 2.1). If this number has reached CoolingIter StoppingFactor (informally, at about the last StoppingFactor temperatures no move has been made-except this analysis can wrap around the border between temperatures), then the algorithm stops and the current seat allocation is its output. 4. CONCLUSIONS For all census years and for both the Banzhaf and Shapley-Shubik power index, the simulated annealing algorithm provides a power balance more in harmony with quotas than do any of the historical algorithms. Of course, this is not overly shocking, as those algorithms were tailored to achieve a certain harmony between quotas and votes apportioned, rather than between quotas and the normalized power vectors induced by the votes apportioned. Generally, the simulated annealing algorithm achieves its strong performance by shifting votes away from the larger state(s), as even the large-state-hostile Adams algorithm is not quite so large-state-hostile as the simulated annealing algorithm. In some sense, the fact (which has been noted elsewhere and which is very apparent in the "State, Quotas, and Normalized Powers" tables of Section 3) that power indices often disproportionally skew power towards very large states-the so-called "big state bias"-is something the simulated annealing algorithm is able to attempt to directly remedy. Thus, if one's notion of fairness is to have a close match between the normalized Banzhaf or Shapley-Shubik power indices of the states and the vector of quotas, the simulated annealing algorithm is a strong contender, and seems to provide a better match than any of the historical algorithms. Our results also suggest that, if one limits one's universe to the six historical algorithms, the United States has chosen the correct one, at least in terms of performance with respect to historical census data: The Huntington-Hill algorithm seems the most power-fair of the historical algorithms, albeit far less fair that simulated annealing. Indeed, in the 1990 tables one can see that simulated annealing and the Huntington-Hill algorithm treated small states (states receiving no more than five seats) identically. The only difference in their apportionments is that, relative to the Huntington-Hill algorithm, simulated annealing shifts voting weight away from large states and towards middle-sized states. The political implications of our results are somewhat surprising. The fact that Adams's algorithm gives large states (e.g., California) far fewer seats than their quotas would seem to entitle them to might suggest that large states are cheated under Adams's algorithm, and indeed both some scholarly analyses [Balinski and Young 1982] and the Supreme Court decision mentioned earlier [Supreme 1992] have been primarily concerned with the correspondence between seats and quotas. Power Balance and Apportionment Algorithms \Delta 15 However, our study shows clearly that the bias power indices show towards large states is so pronounced that it overwhelms the "vote skewing towards small states" of the Adams algorithm. Of course, as the tables show, a large-state-skewed algorithm such as Jefferson's even more strikingly grants disproportionately much power to large states. Overall, in terms of fairness, our results suggest that the current apportionment algorithm used in the United States (the Huntington-Hill algorithm) gives disproportionately much power to large states at the cost of giving disproportionately little power to some middle-sized states, while treating small states fairly. As a final note, nothing above is meant to suggest that the simulated annealing algorithm might be politically feasible. It is quite possible that, in a probabilistic implementation, the vote allocation chosen might depend on the seed given to the random number generator. Worse still it is possible, at least in artificial exam- ples, that different vote allocations would achieve the same degree of quota/power harmony. Thus, the algorithm itself would suggest no preference between the allo- cations, but legislators might well have strong (and differing) preferences. Acknowledgments : The first author was a Bridging Fellow at the University of Rochester's Department of Political Science when this work was started. He thanks that department's faculty for encouraging his interest in political science and voting systems, and for helpful conversations, in particular including suggesting the discussion of the final paragraph of the conclusions section. He also thanks the University of Rochester for the fellowship, and gratefully acknowledges a lifelong debt to Michel Balinski for, many years ago, introducing him both to the study of apportionment and to research. The authors are grateful to William Lucas of the Claremont Graduate School and Peter van Emde Boas of the University of Amsterdam for proofreading an earlier version and for helpful comments. We also thank Joe Malkevitch for helpful comments. We are grateful to ACM Journal on Experimental Algorithmics editor Bernard Moret and two anonymous referees for valuable comments and suggestions. The authors alone, of course, are responsible for any errors. --R Simulated annealing and Boltzmann machines: A stochastic approach to combinatorial optimization and neural computing. Fair Representation: Meeting the Ideal of One Man Fair representation: Meeting the ideal of one man Probabilistic polynomial time is closed under parity reductions. Plan de constitution Mathematical properties of the Banzhaf power index. Mathematics of Operations Research 4 Computers and Intractability: A Guide to the Theory of NP-Completeness Closure properties and witness reduction. Values of large games Values of large games Equations of state calculations by fast computing machines. On closure properties of bounded two-sided error complexity classes An Introduction to Positive Political Theory. Measurement of power in political systems. A method of evaluating the distribution of power in a committee system. The history and status of the P versus NP question. Probability models for power indices. PP is as hard as the polynomial-time hierarchy Counting classes are at least as hard as the polynomial-time hierarchy The complexity of computing the permanent. The complexity of enumeration and reliability problems. --TR

simulated annealing;power indices;apportionment algorithms

297115

Coordination of heterogeneous distributed cooperative constraint solving.

In this paper we argue for an alternative way of designing cooperative constraint solver systems using a control-oriented coordination language. The idea is to take advantage of the coordination features of MANIFOLD for improving the constraint solver collaboration language of BALI. We demonstrate the validity of our ideas by presenting the advantages of such a realization and its (practical as well as conceptual) improvements of constraint solving. We are convinced that cooperative constraint solving is intrinsically linked to coordination, and that coordination languages, and MANIFOLD in particular, open new horizons for systems like BALI.

INTRODUCTION The need for constraint solver collaboration is widely rec- ognized. The general approach consists of making several solvers cooperate in order to process constraints that could not be solved (at least not efficiently) by a single solver. BALI [21, 23, 22] is a realization of such a system, in terms of a language for constraint solver collaboration and a language for constraint programming. Solver collaboration is a glass-box mechanism which enables one to link black-box tools, i.e., the solvers. BALI allows one to build solver collaborations (solver cooperation [25] and solver combination [17]) by composing component solvers using collaboration primitives (implementing, e.g., sequential, concurrent, and parallel collaboration schemes) and control primitives (such as iterators, fixed-points, and conditionals). On the other hand, the concept of coordinating a number of activities, such that they can run concurrently in a parallel and distributed fashion, has recently received wide attention [4, 5]. The IWIM model [1, 2] (Ideal Worker Ideal Manager) is based on a complete symmetry between and decoupling of producers and consumers, as well as a clear distinction between the computational and the coordina- tion/communication work performed by each process. A direct realization of IWIM in terms of a concrete coordination language, namely MANIFOLD [3], already exists. Due to lack of explicit coordination concepts and con- structs, the implementation of BALI does not fully realize its formal model: the treatment of disjunctions and the search are jeopardized and this is not completely satisfactory from a constraint solving point of view. This is mainly due to two causes: (1) the dynamic aspect of the formal model of BALI, and (2) the use of heterogeneous solvers, i.e. , solvers written in different programming languages, with different data rep- resentations. Only a coordination language able to deal with dynamic processes and channels (creation, duplication, dis- /re-/connection), and able to handle external heterogeneous solvers (routines for automatic data conversions) can fullfil the requirements of the formal model of BALI and overcome the problem of its current implementation. This guided us through the different coordination models and lead us to the IWIM model, and the MANIFOLD language. Coordination and cooperative constraint solving are intrinsically linked. This motivated our investigation of a new organizational model for BALI based on MANIFOLD. The results show a wider-than-expected range of implications. Not only the system can be improved in terms of robust- ness, stability, and required resources, but the constraint solving activity itself is also improved through the resulting clarity of search, efficient handling of the disjunctions, and modularity. The system can be implemented closer to its formal model and can be split up into three parts: (1) a constraint programming activity, (2) a solver collaboration lan- guage, and (3) a coordination/communication component. We qualified (and roughly quantified) the improvements co-ordination languages, and more specifically MANIFOLD, can bring to cooperative constraint solving. The conclusions are promising and we feel confident to undertake a future implementation of BALI using MANIFOLD. The rest of this paper is organized as follows. The next section is a brief overview of BALI, its organizational model, and the weaknesses of its implementation. In Section 3, after an overview of MANIFOLD, we describe the coor- dination/communication of BALI using the features of the MANIFOLD system. We then highlight the improvements that we feel are most significant for constraint solving (Sec- tion 4). Finally, we conclude in Section 5 and discuss some future work. BALI [21] is an environment for solver collaboration (i.e., solver cooperation [25, 14] and solver combination [26, 29]) that separates constraint programming (the host language) from constraint solving (the solver collaboration language). The host language is a constraint programming language [34] or possibly a constraint logic programming language [16, 11] which, when necessary, expresses the required solver collaboration through the solver collaboration language. The solver collaboration language supports three strategies called solving strategies. The first strategy consists of determining the satisfiability of the constraint store each time a new constraint occurs ("incremental use of a solver"). The second strategy is an alternative to this method that solves the constraint store when a final state is reached (e.g., the end of resolution for logic programming). The last strategy allows the user to trigger the solvers on demand, for example, to test the satisfiability of the store after several constraints have been settled. Furthermore, BALI allows several solver collaborations, in conjunction with different solving strate- gies, to coexist in a single system. For example, solver S1 can be used incrementally while S2 only executes at the end, and S3 and S4 are always triggered by the user. Since the constraint programming part of BALI is less interesting from the point of view of coordination 1 , this paper focuses on its constraint solving techniques, i.e., the constraint solver collaboration language of BALI. This domain independent language has been designed for realizing a solving mechanism in terms of solver collaborations following certain solving strategies. The basic objects handled by the language are heterogeneous solvers. They are used inside collaboration primitives that integrate several paradigms (such as sequentiality, parallelism, and concur- rency) commonly used in solver combination or cooperation. In order to write finer strategies, we have also introduced some control primitives (such as iterator, fixed-point, and conditional) in the collaboration language. At the implementation level, BALI is a distributed co-operative constraint programming system, composed of a language for solver collaboration (whose implementation allows one to realize servers to which potential clients can connect) plus a host language (whose implementation is a special client of the server). Solver collaboration is a glass-box mechanism which enables one to link black-box tools, i.e., the solvers. Some applications have already used BALI [23]. For exam- ple, a simulation of CoSAc [25] has been realized, and some other solver collaborations have been designed for non-linear constraints. 2.1 The Constraint Solver Collaboration Language Of BALI A detailed description of the solver collaboration language of BALI can be found in [23, 21]. In this section, we give a brief overview of some of the collaboration primitives of BALI. The complete syntax of the solver collaboration language of BALI is given in Figure 1. Sequentiality (denoted by seq) means that the solver E2 will execute on the constraint store C 0 , which is the result 1 The constraint programming part of BALI is described in [21] and [22]. of the application of the solver E1 on the constraint store C. When several solvers are working in parallel (denoted by split), the constraint store C is sent to each and every one of them. Then, the results of all solvers are gathered together in order to constitute a new constraint store analogous to C. Concurrency (denoted by dc) is interesting when several solvers based on different methods can be applied to non-disjoint parts of the constraint store. The result of such a collaboration is the result of a single solver S composed with the constraints that S did not manipulate. The result of S must also satisfy a given property / which is a concurrency function (the set \Psi in Figure 1). For example, basic is a standard function in \Psi that returns the result of the first solver that finishes executing. Some more complex / functions can be considered, such as solved form which selects the result of the first solver whose solution is in a specific solved form on the computation domain. The results of the other solvers (which may even be stopped as soon as S is chosen) are not taken into account. The concurrency primitive is similar to a "don't care'' commitment but also provides control for choosing the new store (using / functions). (positive integers) (boolean observation functions) E) j E) Ar Figure 1: Syntax of the solver collaboration language of BALI These primitives (which comprise the computation part of the collaboration language) can be connected with combinators (which compose the control part, using primitives such as iterators, conditionals, and fixed-points) in order to design more complex solver collaborations. The fixed-point combinator (denoted by f p) repeatedly applies a solver collaboration until no more information can be extracted from the constraint store. This combinator allows one to create an idempotent solver/collaboration from a non-idempotent solver/collaboration. The above primitives and combinators are completely statically defined. We now introduce observation functions of the constraint store which allow one to get more dynamic primitives. These functions are evaluated at run-time (when entering a primitive) using the current constraint store. These functions may be either arithmetic (the set OA in Figure or Boolean (the set OB in Figure 1). Arithmetic observation functions have the profile: Stores ! N . Three such functions are: (1) card var computes the number of distinct variables in the constraint store. This is interesting for solvers that are sensitive to the number of vari- ables. (2) card c returns the number of atomic constraints that comprise the store. This is important for solvers whose complexity is a function of the number of constraints (such as solvers based on propagation). (3) card uni var returns the number of univariate atomic constraints. This is essential for solvers whose efficiency is improved with univariate constraints (such as interval propagation solvers). Boolean observation functions have the profile: Stores ! Boolean. Three such functions are: (1) linear tests whether there exists any variable that occurs more than once in an atomic constraint. This is of interest in deciding the applicability of a linear solver. (2) uni var tests whether there is at least one univariate equality in the store. This information is important since, for example, univariate constraints are generally the starting point of interval propagation. (3) tri tests whether the store is in triangular form (i.e., there are some equality constraints over a variable X, some over variables X and Y , some over X; Y and Z, . This is interesting for eliminating variables, or determining an ordering for the Grobner bases computation. The repeat combinator (denoted by rep) is similar to the fixed-point combinator, but allows applying a solver n times: n is the result of the application of an observation function (or a composition of observation functions) to the constraint store. Since this primitive takes into account the constraint and its form at run-time, it improves the dynamic aspect of the collaboration language. Finally, the conditional combinator (denoted by if) applies one solver/collaboration or another, depending on the evaluation of a condition (which can also depend on observation functions of the constraint store). The following example illustrates the solver collaboration language: seq(A,dc(basic,B,C,D),split(E,F),f p(G)) Consider applying this collaboration scheme to the constraint store c 2 . First A is applied to c and returns c1 . Then, B, C, and D are applied to c1 . The first one that finishes gives the new constraint store c2 . Then E, and F execute on c2 . The solution c3 is a composition of c 0 3 (the solution of E) and c 00 3 (the solution of F). Finally, G is repeatedly applied to c3 until a fix-point, c4 , is reached, which is the final solution of the collaboration. 2.2 Organizational Model And Implementation The role of the organizational model we have implemented is: 1) to create a distributed environment for integrating heterogeneous solvers 3 , 2) to establish communication between solvers in spite of their differences, 3) to coordinate their ex- ecutions. Such an organizational model turns solver collaborations into servers to which clients (such as the implementation of the host language or all kinds of processes requiring a solver) can connect. This model enabled us to implement BALI and create/execute solver collaborations [21]. 2 In order to simplify the explanation, we consider here solvers that return only one solution (one disjunct). We detail the treatment of disjunctions in the next sections. 3 Each solver (software, library of tools, client/server architecture) has its own data representation, is written in a different programming language, and executes on a different architecture and operating system. 2.2.1 Agent The realizations of solvers and solver collaborations are het- erogeneous. However, by an encapsulation mechanism we homogenize the system, and obtain what we call agents. Each agent is autonomous and is created, works, and terminates independently from the others. Hence, agents can execute in parallel or concurrently in a distributed architecture Solvers are encapsulated to create simple agents. As shown in [21], a solver collaboration is a solver. Applying this concept to the architecture, encapsulation becomes a hierarchical operation. Hence, several simple agents can be encapsulated in order to build a complex agent. However, viewed from the outside of a capsule, simple and complex agents are identical. ADMISSIBILITY CONVERTER RECOMPOSITION ECLiPSe -> S INTERNAL ECLiPSe CONVERTER Figure 2: Simple agent In the current implementation of BALI, solvers are encapsulated into ECLiPSe 4 processes (see Figure 2). Hence, ECLiPSe launches the solvers and re-connects their input and output through pipes. The data structure converters are written in Prolog and the data exchanges between capsules and solvers are performed via strings. The encapsulation also provides a constraint store for the solver it represents (a local database for storing the information), an admissibility function (which is able to recognize which constraints of the store can be handled by the solver), and a recomposition function (which recreates an equivalent store using the constraints treated by the solvers, and the constraints not admissible by the solver). The interface of an agent is an ECLiPSe process. Moreover, Prolog terms can be transmitted between two ECLiPSe processes. Inter-agent communication is thus realized with high level terms, and not strings or bits. Furthermore, there is no need for syntactic analyzers between pairs of agents. A complex agent (encapsulation of a solver collabora- behaves like a simple agent, though its internal environment is a bit different (see Figure 3). It has a constraint store for keeping the information it receives: this is its knowledge base. For managing this base, it has a recomposition function which re-builds the constraint store when some agents send some of their solutions. The major work of a complex agent is the coordination (as determined by the collaboration primitive it represents) of the agents it encapsulates. 4 ECLiPSe [20] is the "Common Logic Programming System" developed at ECRC. ECLiPSe COORDINATOR ECLiPSe ENCAPSULATION INTERNAL ECLiPSe RECOMPOSITION ENCAPSULATION Figure 3: Complex agent 2.2.2 Coordination We now describe the coordination of the implementation of BALI (see [21] for more details), but not the coordination of its formal model. An agent can be in one of three different states: running (R), sleeping (S), or waiting (W). When an agent receive a constraint c, it becomes running to solve c. An agent is in the W state when it is waiting for the answer from one or more agents. An agent is in the S state when it is neither running nor waiting. These states together with the communication among agents, enable us to describe the coordination of the constraint solvers. Sequential primitive: seq(S1 ,S2 ,. ,Sn) tries to solve a constraint by sequentially applying several solvers. It first sends a constraint to S1 and waits for a solution c1 for it. When it receives a solution from S1 , it sends it to S2 , waits for a solution c2 , sends it to S3 , and so on, until it reaches Sn . Finally, the solution cn from Sn is forwarded to the superior agent as one of the solutions of the sequential primitive. Since we consider solvers that enumerate their solutions (i.e., each solution represents a disjunct of the complete solution), the sequential agent must wait for the other disjuncts of Sn which will be treated the same way as cn . Backtracking is then performed on Sn\Gamma1 , Sn\Gamma2 and back to S1 . In a sequential collaboration, several agents are "pipelined" and work in "parallel", but the solutions are passed "sequentially" from one agent to the next. primitive: split(S1 ,S2 ,. ,Sn ) applies several solvers in parallel on the same constraints. The solution of split is a Cartesian-product-like re-composition of all the solutions of S1 ,S2 ,. ,Sn . When a split agent receives a solve request from its superior, it forwards it to all its S i 's. Then, it waits and stores all the solutions of each S i . Finally, the split agent creates all the elements of the Cartesian-product of the solutions, and sends them one by one to its superior agent. don't care primitive: dc(/1 ,S1 ,S2 ,. ,Sn) introduces concurrency among solvers. Upon receiving a constraint c from its superior, the don't care agent forwards c to all its sub-agents, S i 's. Then it waits for a solution c 0 from any of its sub-agents. If c 0 does not satisfy /1 5 then c 0 is forgotten and the don't care agent waits for a solution from another sub-agent (other than the one that produced c 0 ). As soon as the don't care agent receives a solution c 0 from some S i 5 /1 is an element of the set / of boolean functions. They test whether or not a constraint satisfies some properties. that satisfies /1 , all other sub-agents are stopped and c 0 , as well as all other solutions produced by S i , are forwarded to the superior agent. fix-point primitive: f p(S) repeatedly applies S on a con- straint, until no more information can be extracted from the constraint. The solving process starts when the fix-point agent receives a constraint c from its superior. It is an iterative process and in each iteration k, we consider a set Ck of disjuncts to be treated by S (e.g., in iteration 1, C1 consists of a single element, c). In iteration k, the mk disjuncts of Ck must be treated by S: the fix-point agent chooses one element of Ck , ck;i , removes it from Ck , sends it to S and collects all the solutions from S. If the 6 solution from S is equal to ck;i (a fix-point has been reached for this dis- junct), the fix-point agent forwards it to its superior agent. Otherwise, the solutions produced by S are added to Ck+1 . The same treatment is applied to all the elements of Ck to complete the set Ck+1 and the solving process enters iteration 1. The process terminates when at the end of iteration k, the set Ck+1 is empty. repeat primitive: The coordination for the repeat primitive is identical to the fix-point collaboration, except that it stops after a given number n of iterations. The number n is computed at run-time: it is the result of the application of the arithmetic function ffi to the current constraint store. The arithmetic function ffi is composed of observation functions of the constraint store (ele- ments in OA, see Figure 1). The solving process starts when the repeat agent receives a constraint c from its superior. First, n is computed: the coordination is analogous to the one of the fix-point primitive. The process terminates at the end of iteration n, when every solution returned by S for every disjunct in Cn is sent to the superior agent. conditional primitive: if(fl,S1,S2) is reather simple. When it receives a constraint c from its superior agent, this primitive applies the function fl to c. The Boolean function is composed of both arithmetic and Boolean observation functions of the constraint store. If fl(c) is true, then c is forwarded to the sub-collaboration S1, otherwise to the sub- collaboration S2. Then, this primtive becomes an intermediary between one of the sub-agents and its superior agent, i.e., as soon as the selected sub-agent sends a solution, it is forwarded immediately to the superior agent. In fact, after evaluation of fl(c) the conditional primitive acts similarly to a sequential primitive having a single sub-agent. 2.3 Weaknesses Of The Implementation Although ECLiPSe provides some functionality for managing processes and communication, it is not a coordination language. Thus, our implementation does not exactly realize the formal model of BALI: some features are jeopardized, or even missing, as described below. Disjunctions of constraints The disjunctions of constraints returned by a solver are treated one after the other, and for some primitives, they are even stored and their treatment is delayed. For the sequential primitive, this does not drastically jeopardize the solving process. But for the fixpoint primitive, this really endangers the resolution. We must wait for all the disjuncts of a given iteration before entering the next one. A solution would be to duplicate the but due to the encapsulation mechanism, this is not 6 When reaching a fix-point, a solver can return only one solution. reasonable. This treatment of disjunction leads to a loss of efficiency, and to a mixed search 7 during solving (which is not completely convenient from the constraint programming point of view). Static architecture Another limitation of BALI is due to the fact that architectures representing collaborations are fixed. Due to some implementation constraints and the limitations of coordination features of ECLiPSe, the collaborations are first completely launched before being used to solve constraints. Thus, we have a loss of dynamics: 1) parts of the architecture are created even when they are not required, 2) agents cannot be duplicated (although this would be interesting for some primitives such as fix-point), and stated before, the disjunctions are not always handled efficiently. Other compromised features Although the formal model of BALI allows the use of "light" solvers, the implementation is not well suited to support such agents: their coarse grain encapsulation uses more memory and CPU than the solver. Thus, mixing heavy solvers (such as GB [10], Maple [12]) and light solvers (such as rewrite rules or transformation rules) is not recommended. No checks are made to ensure that an architecture and its communication channels have been created properly. Management of resources and load balancing are static: before launching a collaboration, the user must decide on which machine the solver will run. 3 MANIFOLD: A NEW COORDINATION FOR BALI We now explain how we can use the coordination language MANIFOLD [3] to significantly improve the implementation of BALI, and remain closer to its formal model. 3.1 The Coordination Language MANIFOLD MANIFOLD is a language for managing complex, dynamically changing interconnections among sets of independent, concurrent, cooperative processes [1]. MANIFOLD is based on the IWIM model of communication [2]. The basic concepts in the IWIM model (thus also in MANIFOLD) are processes, events, ports, and channels. Its advantages over the Targeted-Send/Receive model (on which object-oriented programming models and tools such as PVM [13], PAR- MACS [15], and MPI [7] are based) are discussed in [1, 27]. A MANIFOLD application consists of a (potentially very large) number of processes running on a network of heterogeneous hosts, some of which may be parallel systems. Processes in the same application may be written in different programming languages. The MANIFOLD system consists of a compiler, a run-time system library, a number of utility programs, libraries of built-in and pre-defined processes, a link file generator called MLINK and a run-time configurator called CONFIG. The system has been ported to several different platforms (e.g., SGI Irix 6.3, SUN 4, Solaris 5.2, IBM SP/1, SP/2, and Linux). MLINK uses the object files produced by the (MANIFOLD and other language) compilers to produce link files and the makefiles needed to compose the executables files for each required platform. At the run time of an ap- plication, CONFIG determines the actual host(s), where the processes (created in the MANIFOLD application) will run. 7 The search strategy is breadth-first for the fix-point and repeat primitives, but depth-first for the sequential and don't care primitives. The library routines that comprise the interface between MANIFOLD and processes written in the other languages (e.g., C), automatically perform the necessary data format conversions when data are routed between various different machines. MANIFOLD has been successfully used in a number of applications, including in parallelization of a real-life, heavy duty Computational Fluid Dynamics algorithm originally written in Fortran77 [8, 9, 18], and implementation of Loosely-Coupled Genetic Algorithms on parallel and distributed platforms [31, 33, 32]. 3.2 BALI In MANIFOLD Although BALI solvers are black-boxes and are heterogeneous, this does not cause any problems for MANIFOLD, because it integrates the solvers as external workers. Thus, communication and coordination can be defined among them in the same way as with normal MANIFOLDagents. MANIFOLD can bring many improvements to BALI such as: robustness: managing the faults in the system is not an easy task with ECLiPSe. ffl portability: MANIFOLD runs on several architectures, and requires only a thread facility and a subset of PVM [13]. ffl modularity: in the current implementation, constraint solving is separated from constraint programming. Using MANIFOLD, we can also split up the coordination part from the solving part. ffl extension of the collaboration language: each primitive will be an independent coordinator. Thus, adding a new primitive will be simplified. ffl additional new features: MANIFOLD provides tools to implement certain functionalities that are not available in the current version of BALI (e.g., choice of the machines, light weight processes, architectures, load etc. In the following, we elaborate only on the most significant of the above points, i.e., the ones that make an intensive use of the MANIFOLD features or the ones that are the most significant for constraint solving. 3.2.1 Lighter Agents denoted S Coordinator Rec. Conv Solver Adm Conv Figure 4: Lighter simple agent The current encapsulation (one ECLiPSe process for each solver/collaboration) is really heavy. MANIFOLD can produce lighter capsules using threads to realize filters and workers. They will replace the computation modules of ECLiPSe. Thus, a simple agent (see Figure 4) can consist of: ffl a coordinator for managing the messages and agents inside the encapsulation. This coordinator is also the in/out gate of the capsule (when communicating with superior agents). ffl a solver, which is the same as in the previous implementation ffl four filters (MANIFOLD workers): the first filters the constraints the solver can handle, the second converts the data into the syntax of the solver, the third converts the solutions of the solver into the global syntax 8 and the last re-composes equivalent solutions based on the solutions of the solver and the constraints it cannot handle. A complex agent (see Figure 5) is now the encapsulation of several simple/complex agents together with some filters. The filters and the coordinator (coordinators are described in Section 3.2.2) are specialized for the collaboration primitive the agent represents. For a split collaboration, only one filter is required: a store manager which collects all the solutions from the sub-agent and incrementally builds the elements of their Cartesian-product (as soon as one element is completed, it is sent to the coordinator). In a / don't care primitive, one filter is required for applying the / function to the constraints. For the sequential primitive, as well as the fix-point, no filters are required. seq Coordinator Coordinator split dc Coordinator Sn Sn Sn Store Man. Coordinator fix-point Figure 5: Lighter complex agent This new kind of encapsulation has several advantages. The global architecture representing a solver collaboration will require less processes than before, and also less mem- ory. This is due to several facts: the use of threads instead of heavy processes, the notion of filters, and the sharing of workers, filters, and solvers between several agents (see Figure 6). The creation of another instance of a solver will depend on the activity of the already running instances. Agents are not black-boxes anymore: they become glass- boxes sharing solvers and filters with other agents. But the main advantage is certainly the following: the coordination is now separated from the filters, encapsulated into individual modules, each of which depends on the specific type of collaboration it implements, and can use all the features of MANIFOLD. Thus, it is possible to arrive at a coordination scheme that respects the formal model of BALI. 8 Global syntax is the syntax used in the filters and between agents. Store Sn Coordinator Coordinator Solver Conv Rec. Conv Coordinator dc Coordinator split Conv Conv Solver Adm F S'n Man. Rec. Figure Shared solvers and filters 3.2.2 Coordinators Using MANIFOLD and the new encapsulation process, it is now possible to overcome the problems inherent in the previous implementation of BALI. Dynamic handling of the solvers Since the coordination features are now separated from the filters and workers, the set up of the distributed architecture and its use are no longer disjoint phases. This means that when a solving request is sent, the collaboration will be built incrementally (agent after agent) and only the necessary components will be created. For example, in a conditional or guarded collab- oration, only the "then" or the "else" sub-collaboration will be launched. If another request is sent to the same collab- oration, the launched components will be re-used, possibly augmented by some newly created components. When a solver/collaboration is requested to solve a con- straint, several cases can arise. If the solver/collaboration S has not already been launched, then an instance of S will be created. If it is already launched but all of its instances are busy (i.e., all instances of S are currently working on con- straints) another instance will be created. Otherwise, one of the instances will be re-used for the new computation. The function find instance manages this functionality (see Appendix A.1). Dynamic handling of the disjunctions Contrary to the current implementation 9 , disjunctions are treated dynami- cally. We demonstrate this for the sequential collaboration seq(S1 ,S2 ,. ,Sn ). All the disjuncts produced by S1 must be sent to S2 . With ECLiPSe, a disjunct c1 of S1 is completely solved by S2 , . , Sn (meaning all possible disjuncts created by S2 , . , Sn are produced), before treating the next disjunct c2 of S1 . MANIFOLD allows us to use pipelines to solve c2 as soon as it is produced by S1 . If S2 is still working on c1 , and all the other instances of S2 are busy, then a new instance of S2 is created for solving c2 . The treatment of c2 is no longer postponed. This mechanism applies to all sub-agents of the sequential agent. This introduces a new problem: there may be a combinatorial explosion of the number of instances of S2 , . , Sn . However, this can rarely happen: while an agent S i is producing solutions, the agent S i+1 is already solving (and has already solved) some of the previous constraints. Thus, 9 Currently, the fix-point coordinator waits for all the solutions of the sub-collaboration before entering the next iteration. some instances have already returned to a sleeping state and can be re-used. Nevertheless, the following case may arise. Suppose the solvers S i are arranged such that as the index grows, the designated solvers, S i , become slower, and suppose every S i creates disjuncts. The number of instances will become exponential in this case, and the system will therefore run out of resources. In order to overcome this problem, the number of instances can be limited (see Appendix A.1). Thus, when a solving request is to be sent to the agent S, and the maximal number of (its) instances is reached, and all its instances are busy, the superior agent will wait for the first instance to return to the sleeping state. This mechanism does not imply a completely dynamic treatment of the disjunctions. However, it gives a good compromise between the delay for solving a disjunct and the physical limitation of the resources. Coordinators for the primitives We now describe the coordinators for the sequential primitive. Some other primitives are detailed in Appendix A. The algorithms are presented here in a Pascal-like language extended with an event functionality. We consider a queue of messages m from p meaning that the message m was received on the port p. task m from p alg means that we remove the message m from the port p and execute the algorithm alg (the message m from p is the condition for executing the task alg). The latter cannot be interrupted. end is a message that is sent by an agent when it has enumerated all its disjuncts. The agents have a number of flags representing the states described in Section 2.2.2. Figure 7: Duplication of a sequential primitive coordinator for seq(S1,.,Sn) S1.Sn: sub-agents; S0: sup-agent ports: p.0.in . p.n.in % for 0=!i!n+1 p.i.in is linked to p.Si.out of Si p.0.1.out . p.0.n+1.out % for 0!i!n+1 p.0.i.out is linked to p.Si.in of Si % p.0.n+1.out is linked to p.S0.in of S0 task c from p.i.in: if j!?NULL then send c to p.j.i+1.out; Mi=Mi+1 else send c to p.i.in % c is sent again and again to p.i.in % till an instance of Si+1 becomes free task end from p.i.in: and Si-1 is Sleeping then Si is Sleeping task end from p.n.in: and Sn-1 is Sleeping then Sn is Sleeping; send end to p.0.n+1.out task c from p.i.in is used to forward a solution from S i to S i+1 . If no instance of S i+1 is free, and it is not possible to create a new instance, then the same message is sent again, and will be treated later. To detect the end of a sequential primitive, we count the solutions and end messages of each of the sub-agent. An agent S i becomes sleeping when S i\Gamma1 is already sleeping, and when S i produces as many end messages as the number of solutions of S i\Gamma1 . The superior agent S0 is never duplicated inside a collab- oration, since the coordinator can create only a sub-architec- ture; the collaboration does not duplicate itself. That is the job of the superior agent: it either finds a free instance of the collaboration, or creates one if the maximal number of instances is not yet reached (see Figures 7 and 8 for an example of duplication). dc dc Y YS SS Figure 8: Duplication of a /-don't care primitive We have seen that coordination languages, and MANIFOLD in particular, are helpful for implementing cooperative constraint solving. However, the advantages are not only at the implementation level. MANIFOLD allows an implementation closer to the formal model of BALI, and this implies some significant benefits for constraint solving: faster execution time, better debugging, and clarity of the search [28] during constraint solving (see Table 9). The architecture also gains through some improvements: robustness, reliabil- ity, quality, and a better management of the resources (see Table 10). This last point also has consequences for the end user: as the architectures representing a solver collaboration become lighter, the end user can build more and more complex collaborations, and thus, solve problems that could not be tackled before. Constraint solving Treatment of disjunctions is a key point in constraint solving. The most commonly required search is depth-first: each time several candidates appear, take one, and continue with it until reaching a solution, then backtrack to try the other candidates. One of the reasons for this choice is that, generally, only one solution is re- quired. Contrary to the first implementation of BALI, the coordination we described with MANIFOLD leads to what we call a "parallel depth-first and quick-first" search. The parallel depth-first search is obvious. The quick-first search arises from the fact that each constraint flows through the agents independently from the others. Hence (ignoring the boundary condition of reaching the instance limits of solvers, mentioned above), it is never delayed by another constraint, nor stops at the input of a solver or in a queue. The result is that the solution which is the fastest to compute (even if it is not originated from the first disjunct of a solver) has a better chance to become the first solution given by the solver collaboration 10 . Debugging, collaboration improvement, and graphical interface to present output will be eased. The coordinators can duplicate the messages and send them to a special worker. This latter can then be linked to a display window (text or graphic) or a profiler. It will enable users observe the flow of data in a collaboration. Thus, users can extract statistics on the utilization of the solvers and draw conclusions on the efficiency of a newly designed collaboration. All this process can lead to a methodology for designing solver collaborations Due to its encapsulation techniques, the current implementation jeopardizes the use of "fine grain" solvers (solvers that require little memory and CPU). Although we can envisage encapsulating a single function with an ECLiPSe pro- cess, this is not reasonable. Though not really designed for fine grain agents, MANIFOLD still gives more freedom to use single functions (such as rewrite rules or constraints trans- formations) as solvers. With MANIFOLD, single functions for simplifying the constraints can easily be inserted in a collaboration as threads without compromising the efficiency of the whole architecture; this significantly enlarges the set of solvers that can be integrated in BALI. BALI in: ECLiPSe MANIFOLD search during solving mixed depth-first first solution - ++++ execution time treatment of the disjunctions - +++ use of "fine grain" solvers - ++ add of solvers (encapsulation) extension of the collaboration language - +++ "debugging tools" - +++ improvement of solver collab. - +++ input graphical interface - ++ output graphical interface Figure 9: Improvements for constraint solving Coordination is now separated from collaboration: a collaboration primitive implies a coordinator separated from the converters, recomposition functions, and admissibility When a branch leads faster to a solution, we find it quickly, because we do not have to explore all branches before this one. functions. Thus, with the same filters we can easily implement new primitives: only the coordinator has to be modi- fied, and in some cases a filter must be added. Architecture The major limitation of BALI is the large amount of resources it requires. Of course, this is an intrinsic problem with cooperative solvers: they are generally costly in memory, CPU, etc. But another limiting factor is the overhead of the current encapsulation mechanism. With the new encapsulation technique, MANIFOLD will decrease the required resources. Furthermore, with dynamic handling of disjunctions, we expect the new architecture to be less voracious. BALI in: ECLiPSe MANIFOLD construction of the architecture static dynamic robustness - +++ extension of BALI stability - ++ graphical interface (in/output) - ++ quality of communication - ++++ coordination functionalities number of processes but solvers - number of communications ++ - Figure 10: Improvements for architectures The system will gain in robustness, since currently no failure detection of the architecture is possible. The collaborations will be more stable and less susceptible to broken communication and memory allocation problems. The dynamic building of the architecture will decrease the number of unnecessary processes: only the agents required in a computation are launched. The only negative point is the increased communica- tion. With the current implementation, the encapsulation is composed of two communicating processes: ECLiPSe and a solver. All the filters are modules of the ECLiPSe process. With MANIFOLD, the filters are independent agents that also exchange information. However, this should not create a bottle-neck since messages are generally short, communicating agents are usually threads in the same process that using shared memory to communicate, and no single agent conducts nor monitors all communication. 5 CONCLUSION We have introduced an alternative approach for designing cooperative constraint solving systems. Coordination lan- guages, and MANIFOLD in particular, exhibit properties that are appropriate for implementing BALI. However, implementation improvement is not the only advantage. Using MANIFOLD we can produce a system closer to the formal model of BALI, and some significant benefits are also obvious for constraint solving. The major improvements are the treatment of the disjunctions, the homogenization of search, and the reduction of required resources. A fare management of the disjuncts returned by a solver often leads to quicker solutions. Moreover, due to replication, the complete set of solutions is always computed more efficiently. Although the mixed search used in the current implementation of BALI does not really influence resolution when looking for all the solutions of a problem, it becomes a real nuisance when looking for only one solution. Furthermore, observing the resolution and following the flow of constraints is not conceiv- able. MANIFOLD overcomes this problem by providing a "parallel depth-first and quick-first" search: each disjunct is handled independently, and thus no constraint resolution is delayed or queued. Comparing BALI to other systems (such as cc [30] and Oz [19]) is not easy since they do not have the same objectives [21]. cc is a formal framework for concurrent constraint programming, and Oz is a concurrent constraint programming system. However, one of the major distinctions is that BALI, contrary to Oz and cc, enables the collaboration of heterogeneous solvers. Another essential difference concerns the separation of tasks. With Oz and cc, constraint pro- gramming, constraint solving, and coordination of agents are mixed. With BALI, constraint programming is separated from cooperative constraint solving, and using MANIFOLD, cooperative constraint solving is split up into coordination of agents and constraint solving: each aspect of cooperative constraint programming is an independent task. Since the implementation model of BALI with MANIFOLD is clearly defined, we can surely start with the implementation phase. Moreover, we know the feasibility of the task, and have already qualified (as well as roughly quantified) the improvements. Hence, we know that it is a worthwhile work. In the future, we plan to integrate a visual interface to assist programmers in writing more complex solver collab- orations. This can be achieved using Visifold [6] and some predefined "graphical" coordinators. In order to perform optimization, we are thinking of adding another search technique to BALI: a best solution search (branch and bound). This kind of search is generally managed by the constraint language. However, MANIFOLD coordinators that represent collaboration primitives can perform the following tasks: they can eliminate the disjuncts that are above the current "best" solution, and also manage the updating of the current best solution. Branching can, thus, be improved and performed sooner. The constraint solver extension mechanism of SoleX [24] consists of rule-based transformations seen as elementary solvers. Until now, the implementation of SoleX with BALI was not really conceivable: rule-based transformations are too fine grain solvers to be encapsulated. With the new model, the implementation of SoleX becomes reasonable. Finally, we are convinced that cooperative constraint solving is intrinsically linked to coordination, and that coordination languages open new horizons for systems like BALI. --R Coordination of massively concurrent activities. The IWIM model for coordination of concurrent activities. Manifold20 reference manual. What do you mean Coordination languages for parallel programming. Visifold: A visual environment for a coordination language. An introduction to the MPI standard. Restructuring sequential Fortran code into a paral- lel/distributed application Using coordination to parallelize sparse-grid methods for 3D CFD problems Maple V user's guide and reference manual. A symbolic-numerical branch and prune algorithm for solving non-linear polynomial sys- tems PARMACS v6. Constraint Logic Pro- gramming: a Survey Combining symbolic constraint solvers on algebraic domains. Multiple semi-coarsened multigrid for 3d cfd ECLiPSe User Manual. Collaboration de solveurs pour la programmation logique 'a contraintes. PhD Thesis, Universit'e Henri Poincar'e-Nancy I An environment for designing/executing constraint solver collaborations. The Constraint Solver Collaboration Language of BALI. SOLEX: a Domain-Independent Scheme for Constraint Solver Ex- tension Implementing Non-Linear Constraints with Cooperative Solvers Simplifications by cooperating decision procedures. search strategies for computer problem solving Cooperation of decision procedures for the satisfiability problem. Concurrent Constraint Programming. Distributed evolutionary optimization in manifold: the rosenbrock's function case study. Parallel and distributed evolutionary computation with Manifold. Parallel evolutionary computation: Multi agents genetic algorithm. Strategic Directions in Constraint Programming. --TR --CTR Monfroy , Carlos Castro, Basic components for constraint solver cooperations, Proceedings of the ACM symposium on Applied computing, March 09-12, 2003, Melbourne, Florida

constraint solver cooperation;dynamic coordination;solver collaboration language;coordination model

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Coordinating autonomous entities with STL.

This paper describes ECM, a new coordination model and STL its corresponding language. STL's power and expressiveness are shown through a distributed implementation of a generic autonomy-based multi-agent system, which is applied to a collective robotics simulation, thus demonstrating the appropriateness of STL for developing a generic coordination platform for autonomous agents.

Introduction Coordination constitutes a major scientific domain of Computer Science. Works coming within Coordination encompass conceptual and methodological issues as well as implementations in order to efficiently help expressing and implementing distributed appli- cations. Autonomous Agents, a discipline of Artificial Intelligence which enjoys a boom since a couple of years, embodies inherent distributed applications. Works coming within Autonomous Agents are intended to capitalize on the co-existence of distributed entities, and autonomy-based Multi-Agent Systems (MAS) are oriented towards interactions, collaborative phenomena and autonomy. Today's state of the art parallel programming mod- els, such as (distributed) shared memory models, data and task parallelism, and parallel object oriented models (for an overview see [30]), are used for implementing general purpose distributed appli- cations. However they suffer from limitations con- Part of this work is financially supported by the Swiss National Foundation for Scientific Research, grants 21-47262.96 and 20-05026.97 cerning a clear separation of the computational part of a parallel application and the "glue" that coordinates the overall distributed program. Especially these limitations make distributed implementations burdensome. To study problems related to coordi- nation, Malone and Crowston [25] introduced a new theory called Coordination Theory aimed at defining such a "glue". The research in this area has focused on the definition of several coordination models and corresponding coordination languages, in order to facilitate the management of distributed applications. Coordination is likely to play a central role in MAS, because such systems are inherently dis- tributed. The importance of coordination can be illustrated through two perspectives. On the one hand, a MAS is built by objective dependencies which refers to the configuration of the system and which should be appropriately described in an implementation. On the other hand, agents have subjective dependencies between them which requires adapted means to program them, often involving high-level notions such as beliefs, goals or plans. This paper presents STL, a new coordination language based on the coordination model ECM, which is a model for multi-grain distributed applications. STL is used so as to provide a coordination frame-work for distributed MAS made up of autonomous agents. It enables to describe the organizational structure or architecture of a MAS. It is conceived as a basis for the generic multi-agent platform CODA 1 . Coordination Theory, Models and Languages Coordination can be defined as the process of managing dependencies between activities [25], or, in the field of Programming Languages, as the process of building programs by gluing together active pieces [10]. To formalize and better describe these interdependencies it is necessary to separate the two essential parts of a parallel application namely, computation and coordination. This sharp distinction is also the 1 Coordination for Distributed Autonomous Agents. key idea of the famous paper of Gelernter and Carriero [10] where the authors propose a strict separation of these two concepts. The main idea is to identify the computation and coordination parts of a distributed application. Because these two parts usually interfere with each other, the semantics of distributed applications is difficult to understand. To fulfill typical coordination tasks a general coordination model in computer science has to be composed of four components (see also [20]): 1. Coordination entities as the processes or agents running in parallel which are subject of coordination 2. A coordination medium: the actual space where coordination takes place; 3. Coordination laws to specify interdependencies between the active entities; and 4. A set of coordination tools. In [10] the authors state that a coordination language is orthogonal to a computation language and forms the linguistic embodiment of a coordination model. Linguistic embodiment means that the language must provide language constructs either in form of library calls or in form of language extensions as a means to materialize the coordination model. Orthogonal to a computation language means that a coordination language extends a given computation language with additional functionalities which facilitate the implementation of distributed applications. The most prominent representative of this class of new languages is Linda [9] which is based on a tuple space abstraction as the underlying coordination model. An application of this model has been realized in Piranha [8] (to mention one of the various applications based on Linda's coordination model) where Linda's tuple space is used for networked based load balancing functionality. The PageSpace [14] effort extends Linda's tuple space onto the World-Wide-Web and Bonita [28] addresses performance issues for the implementation of Linda's in and out primitives. Other models and languages are based on control-oriented approaches (IWIM/Manifold [2] [3], ConCoord [18], Darwin [24], ToolBus [5]), message passing paradigms (CoLa [17], Actors [1]), object-oriented techniques (Objective Linda [19], JavaSpace [31]), multi-set rewriting schemes (Bauhaus Linda [11], Gamma [4]) or Linear Logic (Linear Objects [6]). A good overview on coordination issues, models and languages can be found in [27]. Our work takes inspiration from control-oriented models and tuple-based abstractions, and focuses on coordination for purpose of MAS distributed implementations 3 Coordination using Encapsu- lation: ECM ECM 2 is a model for coordination of multi-grain distributed applications. It uses an encapsulation mechanism as its primary abstraction (blops), offering structured separate name spaces which can be hierarchically organized. Within these blops active entities communicate anonymously through connections, established by the matching of the entities' communication interfaces. ECM consists of five building blocks: 1. Processes, as a representation of active entities; 2. Blops, as an abstraction and modularization mechanism for group of processes and ports; 3. Ports, as the interface of processes/blops to the external world; 4. Events, a mechanism to react to dynamic state changes inside a blop; 5. Connections, as a representation of connected ports. Figure 1 gives a first overview of the programming metaphor used in ECM. According to the general characteristics of what makes up a coordination model and corresponding coordination language, these elements are classified in the following way: 1. The Coordination Entities of ECM are the processes of the distributed application; 2. There are two types of Coordination Media in ECM: events, ports, and connections which enable coordination, and blops, the repository in which coordination takes place; 3. The Coordination Laws are defined through the semantics of the Coordination Tools (the operations defined in the computation language which work on the port abstraction) and the semantics of the interactions with the coordination media by means of events. An application written using the ECM methodology consists of a hierarchy of blops in which several processes run. Processes communicate and coordinate themselves via events and connections. Ports serve as the communication endpoints for connections which result in pairs of matched ports. Encapsulation Coordination Model. Connection Blop Process Event Figure 1: The Coordination Model of ECM. 3.1 A blop is a mechanism to encapsulate a set of ob- jects. Objects residing in a blop are per default only visible within their "home" blop. Blops are an abstraction for an agglomeration of objects to be coordinated and serve as a separate name space for port objects, processes, and subordinated blops as well as an encapsulation mechanism for events. In Figure 1, two blops are shown. Blops have the same interface as processes, i.e. a name and a possibly empty set of ports, and can be hierarchically structured. 3.2 Processes A process in ECM is a typed object, it has a name and a possibly empty set of ports. Processes in the ECM model do not know any kind of process iden- tification, instead a black box process model is used. A process does not have to care about to which process information will be transmitted or received from. Process creation and termination is not part of the ECM model and is to be specified in the instance of the model. 3.3 Ports Ports are the interface of processes and blops to establish connections to other processes/blops, i.e. communication in ECM is handled via a connection and therefore over ports. Ports have names and a set of well defined features describing the port's charac- teristics. Names and features of a port are referred to as the port's signature. The combination of port features results in a port type. Features. Ports are characterized through a set of features from which the communication feature is mandatory and must be supported by all ECM re- alizations. The communication feature materializes the communication paradigm: it includes point-to- point stream communication (with classical message-passing semantics), closed group (with broadcast se- mantics) and blackboard communication. Additional port features specify e.g. the amount of other ports a port may connect to, see STL as an example. Matching. The matching of ports is defined as a relation between port signatures. Four general conditions must be fulfilled for two ports to get matched: (1) both have compatible values of fea- tures; (2) both have the same name; (3) both belong to the same level of abstraction, i.e., are visible within the same hierarchy of blops; and (4) both belong to different objects (process or blop). Conceptually the matching of process ports can be described as follows. When a process is created in a blop, it creates with its port signature a "potential" in the blop where it is currently embedded. If conditions are fulfilled for two potentials in a blop, the connection between their corresponding ports is established and the potentials disappear. 3.4 Connections The matching of ports results in the following connections ffl Point-to-point Stream. 1:1, 1:n, n:1 and n:m communication patterns are possible; ffl Group. Messages are broadcast to all members of the group. A closed group semantics is used, i.e. processes must be member of the group in order to distribute information in it; ffl Blackboard. Messages are placed on a blackboard used by several processes; they are persistent and can be retrieved more than once in a sequence defined by the processes. 3.5 Events Events can be attached to conditions on ports of blops or processes. These conditions will determine when the event will be triggered in the blop. Condition blop world - // Process definition process p1 - P2Pin port1 !"INPUT"?; BB port2 !"BB"?; // with its ports . // More processes . // A new blop . // More blops create process p1; // Create processes create blop b1(); // Create blops - // End of blop world Figure 2: Layout of a typical program written in STL. checking is implementation dependent (see STL's event definition as an example of how to define event semantics on ports). 4 The Coordination Language We designed and implemented a first language binding of the ECM model, called . STL is a realization of the ECM model applied to multi-threaded applications on a LAN of UNIX workstations. STL materializes the separation of concern as it uses a separate language exclusively reserved for coordination purposes and provides primitives which are used in a computation language to interact with the entities. The implementation of STL [21] is based on Pt-pvm [22], a library providing message passing and process management facilities at thread and process level for a cluster of workstations. In particular, blops are represented by heavy-weight UNIX processes, and ECM processes are implemented as light-weight processes (threads). The ECM model is realized in an STL program whose general structure is outlined in Figure 2. Starting from the default blop world, a hierarchy of processes and blops can be defined, showing the hierarchical structure of the language at definition level. 4.1 STL's Specialities In this Section we look at particularities for the instantiation of the ECM model in STL. 3 Simple Thread Language. Blops The name of a blop is used to create instances of a blop object. Blop objects can be placed onto a specific physical machine, or can be distributed onto a cluster of workstations. The creation of a blop is a complex recursive procedure: it includes the initialization of all static processes and ports defined for this blop and subordinated blops. Figure 3 shows the definition of two blops (called world and sieve) in STL syntax, the line 'create blop sieve s()' initializes the blop somewhere on the parallel machine. The statement could be annotated with a machine name to specify the actual workstation on which the blop should be initialized. The port definitions will be explained later. blop world - blop sieve - // Declaration blop sieve // Two ports: types and names Group a !"CONNECTOR"?; create blop sieve s(); // Create blop Figure 3: Blop declaration and invocation in STL. ECM processes in STL can be activated from within the coordination language and in the computation language. In the coordination language this is done through the instantiation of a process object inside a blop. To dynamically create new processes the process object instantiation can be done in the body of an event or in the computation language directly. To some extent this is a trade-off regarding our goal to totally separate coordination and computation at code level. However, in order to preserve a high level of flexibility at application level, we allow these two possibilities. Process termination is implicit: once the function which implements the process inside the computation language has terminated, the process disappears from the blop. Figure 4 shows an example of an STL process type worker with two static ports in and res and a thread entry point worker; the syntax and semantics of the port definitions will be explained in the next section. Ports and Connections knows static ports as an interface of a process or blop defined in the coordination language, and dynamic ports which are created at runtime in the computation language. However, the type of a dynamic Attribute P2P BB Group MyPort Explanation Communication stream blackboard group stream Communication structure Saturation 1 * 5 Amount of ports that may connect Capacity * 12 Capacity port: in data items Synchronization async async async async Semantics message passing model Orientation inout inout inout in Direction of data flow Table 1: STL's built-in ports and a user defined port with corresponding port attribute values. Process type worker process worker - P2Pin in !"WORK"? // input port P2Pout res !"RESULT"? // output port create process worker w;// An instance Figure 4: Process declaration and invocation in STL. port, i.e, its features must be determined in the co-ordination language. ports use a set of attributes as an implementation for ECM's port features. These attributes must be compatible in order to establish a connection between two ports. Table 1 gives an overview of the attributes of a port; combinations of attributes lead to port types. provides the following built-in port types: point-to-point output ports, (P2Pout), point-to-point input ports (P2Pin), point-to-point bi-directional ports (P2P), groups (Group) and blackboards (BB). Variants of these types are possible and can be defined by the user. The classical stream port. Two matched ports of this type result in a stream connection with the following semantics: Every send operation on such a port is non blocking, the port has an infinite storage capacity (in STL, infinity is symbolized by *), and matches to exactly one other port. The orientation attribute defines whether the port is an output port (P2Pout), an input port (P2Pin), or both (P2P). Group: A set of Group ports form the group mechanism of STL. Ports of this type are gathered in a group and all message send operations are based on broadcast, that is, the message items will always be transfered to all members of the group. A closed group semantics is used. BB: The BB stands for Blackboard and the resulting connection has a blackboard semantics as defined for ECM. In contrast to the previous port types, messages on the blackboard are now persistent objects and processes retrieve messages using a symbolic name and a tag. Combinations of these basic port types are possi- ble, for example to define a (1:n) point-to-point con- nection, the saturation attribute of a P2P port can be augmented to n, see Table 1 port MyPort. Synchronous communication can be achieved by changing the type of message synchronization to syn- chronous, thus yielding in point-to-point synchronous communication. For 1:n this means that the data producing process blocks until all the n processes have connected to the port, and every send operation returns only after all n processes have received the data item. In STL the synchronization attribute overrides the capacity attribute, because synchronous communication implies a capacity of zero. However, asynchronous communication can be made a little bit less asynchronous by setting the capacity attribute to a value n to make sure that a process blocks after having sent n messages. Note that, the capacity attribute is a local relation between the process and its port: for asynchronous communication with a certain port capacity, it is only guaranteed that the message has been placed into the connection which does not necessarily mean that another process connected via a port to this connection has (or will) actually received the message. Connections result in matched ports and are defined in accordance to the ECM model. STL's Events Events are triggered using a condition operation on a port. The event is handled by an event handler inside the blop. Conditions related to ports of processes or blops determine when the event will be executed in the blop (for an overview on port conditions, see Table 2). Whether an event must be triggered or not will be checked by the system every time data flows through it or a process accesses it. Otherwise a condition like isempty would uninterruptedly trigger events for ports of processes, because at start-up of the process ports are empty. Condition Description unbound(port p) For saturation 6= * the predicate unbound() returns true if the port has not yet matched to all its potential communication partners. For ports with saturation=*, the unbound predicate returns always true. accessed(port p) Equals true whenever the port has been accessed in general. isempty(port p) Checks whether the port has messages stored or not. isfull(port p) Returns true if the port's capacity has been reached. msg handled(port p, int n), less msg handled(port p, int n) Equals true if n messages, or less than n messages have been handled, respectively. Table 2: Conditions on ports. event new-worker() - create process worker new; when unbound(new.out) then new-worker(); Process type worker process worker - P2Pin in !"WORK"? // in port P2Pout out !"WORK"? // out port create process worker w; // Attach event to port when unbound(w.out) then new-worker(); Figure 5: An example of event handling in STL. After an event has been triggered, a blop is not tuned anymore to handle subsequent events of the same type. In order to handle these events again, the event handling routine must be re-installed which is usually done in the event handling routine of the event currently processed. Very useful is the unbound condition on ports because it enables the construction of parallel software pipelines very elegantly. If we reconsider Figure 4 and extend it to Figure 5, we see the interaction of event conditions and ports in STL. First an event new worker is declared. The event is attached to an unbound condition on the out port of the initial process denoted w.out. If process w either reads or writes data from/to its port out, the event new worker is triggered because at that time there are no other ports to which w.out is currently bound, so unbound(w.out) returns TRUE. The event body of the event declaration of new worker creates a new process of type worker, resulting in a new port signature or "potential" in the blop. The blop matches now the two ports (w.out and new.in) and the information can be transferred from w to new. The same mechanism recursively works for the out port of the created process new, because the same condition (unbound) is attached to its port new.out. This example illustrates the necessity of reinstalling explicitly event handlers so as to ensure a coordinated execution of the event handler body; in this case a process creation. Primitives STL is a separate language used in addition to a given computation language (in this case C), however the coordination entities must be accessed from within the computation language. Therefore, we implemented a C library to interact with the coordination facilities of STL. The set of primitives includes operations for creating dynamic ports, methods to transfer data from and to ports, and operations for process management; see [21] for a detailed specification 4.2 STL Compiler and Runtime Syste Figure 6 shows the basic building blocks of the STL programming environment in context with Pt-pvm. A distributed application consists of two files: the co-ordination part (app.stl) and the computation part (app-func.c). The STL part will be parsed by the compiler to produce pure Pt-pvm code. The final program will then be linked with the runtime libraries of STL and Pt-pvm, the user supplied code, and the generated code of the STL compiler. app-func.o app.stl app.c app-func.c app.o stllib.a User User Compiler Compiler Linker app ptpvm.a Figure Programming environment of STL. 5 Coordination of Autonomous Agents in STL One of our target with STL is the distributed implementation of a class of multi-agent systems. The multi-agent methodology is a recent area of distributed artificial intelligence. A MAS is an organization given by a set of artificial entities acting in an environment. Focusing on collective behaviors, this methodology is aimed at studying and taking advantage of various forms of agent influences and interactions. It is widely used either in pure simulation of interacting entities (for instance in artificial life [23]) or in problem solving [26]. Definitions of MASs are general enough to address multiple domains which can be specified by the nature of the agents and the environment. We will proceed by composing a typical software MAS starting from problematics logically inducing a "dis- tributed" approach and which could capitalize on self-organizing collective phenomena. One key concept in such approaches is emergence, that is the apparition of functional features at the level of the system as a whole. Our work is aimed at leading to robust solutions for applications in the frameworks of robotics and parallelism. For the design of our sys- tems, we follow the "new AI" trend [7], [13]: our agents are embodied and situated into an environ- ment. The primeval feature we attempt to embody into them is autonomy: the latter is believed to be a necessary condition for flexibility, scalability, adaptability and emergence. In what follows, we will formally describe the class of MASs we attempt to implement on distributed ar- chitectures. Then we will motivate our project and discuss the difficulties encountered when distributing such MASs. We will eventually present the implementation with STL of a peculiar application belonging to our class of MASs. 5.1 A Generic Model for an Autonomous Agents' System Our generic model is composed of an Environment and a set of Agents. The Environment encompasses a list of Cells, each one encapsulating a list of on-cell available Objects at a given time (objects to be manipulated by agents) and a list of connections with other cells, namely a Neighborhood which implicitly sets the topology. This way of encoding the environment allows the user to cope with any type of topol- ogy, be it regular or not, since the set of neighbors can be specified separately for every cell. The general architecture of an agent is displayed on Figure 7. An agent possesses some sensors to perceive the world within which it moves, and some effectors to act in this world (embodiment). The implementation of the different modules presented on Figure 7, namely Perception, State, Actions and Control algorithm depends on the application and is under the user's responsibility. In order to reflect embodiment and situatedness, perception must be local: the agent perceives only the features of one cell, or a small sub-set of cells, at a given time. The control algorithm module is particularly important because it defines the type of autonomy of the agent: it is precisely inside this module that the designer decides whether to implement an operational or a behavioral autonomy [33]. Operational autonomy is defined as the capacity to operate without human intervention, without being remotely controlled. Behavioral autonomy supposes that the basis of self-steering originates in the agent's own capacity to form and adapt its principles of behavior: an agent, to be behaviorally au- tonomous, needs the freedom to have formed (learned or decided) its principles of behavior on its own (from its experience), at least in part. For instance, a very basic autonomy would consist of randomly choosing the type of action to take, a more sophisticated one would consist of implementing some learning capabil- ities, e.g. by using an adaptive neural network. 5.2 A Typical Application: Gathering Agents Our class of MAS can support numerous applica- tions. We illustrate with a simulation in the frame-work of mobile collective robotics. Agents (an agent simulates the behavior of a real robot) seek for objects distributed in their environment, and we would like them to stack all objects, like displayed in Figure 8. Our approach rests on a system integrating operationally autonomous agents, that is, each agent in the system acts freely on a cell (the agent decides which action to take according to its own control algorithm and local perception). Therefore, there is no master responsible for supervising the agents in the system, thus allowing it to be more flexible and fault tolerant. Agents have neither explicit coordination features for Sensors Effectors Content and connections of occupied Cell Perception Actions objects, move Manipulate algorithm Control State Figure 7: Architecture of an agent. Figure 8: Collective robotics application: stacking objects detecting and managing antagonism situations (con- flicts in their respective goals) nor communication tools for negotiation. In fact they communicate in an indirect way, that is, via their influences on the environment. In our simulation, the environment is composed of a discrete two dimensional L cell-sided grid, a set of N objects and n mobile agents. An object can be located on a cell or carried by an agent. Under the given constraints, we implemented several variants for agent modules (details can be found in [12] and [16]). A simple control algorithm that can be used is as follows. Agents move randomly from a cell to a connected one. If an agent that does not carry an object comes to a cell containing NO objects, it will pick up one object with a probability given by 1=NO ff , where ff - 0 is a constant; if an agent that carries an object comes to a cell containing some objects, it will systematically drop its object. If the cell is empty, nothing happens. This simulation has been already serially imple- mented, exhibiting the emergence of properties in the system, such as cooperation yielded by the recurrent interactions of the agents; agents cooperate to achieve a task without being aware of that. Further details about this simulation and outcomes can be found in [16]. 5.3 Coexistence and Distribution of Autonomous Agents We would like to stress on the problem of implementing a MAS, such as the one described above, on a distributed architecture. The instantiation of a MAS is most of the time se- rial. However that may be, because one draws inspiration from group of robots or living entities, it seems clear that the agents may run in parallel. At the level of abstraction of the MAS specification, the term parallelism simply conveys the notion co-existence of the autonomous agents. Hence, parallelism ideally underlies every conception of MAS and is thought to be implicitly taken into account in a serial imple- mentation. Paradoxically, the projection of a MAS onto a distributed architecture turns out to be far from being obvious. We will not discuss here what are the fundamental origins of this difficulty; more details may be found in [15] (it concerns the different concepts of time in MAS and in distributed comput- ing, respectively). We will illustrate the difference between concepts of coordination in both distributed computing and MAS. The problem crystallizes around the different levels of what we understand when speaking about coordination In a conception stage (i.e. before any implemen- tation), the notion of "coordination of agents" refers to a level of organization quite different from the one at programming level. We studied our MAS application (the gathering agents) in order to investigate some methods of cooperation between agents, namely emergent or self-organized cooperation, a sub-domain of coordination. In this stage, "cooperation between agents" deals with dependencies between agents as autonomous entities. The challenge is to find an appropriate trade-off between cooperation, the necessary fruitful coordination of inter-dependent entities, and their relative autonomy. At this stage, the envisaged cooperation methods must be described, in terms of the agents' architecture, their perception, their actual sensors and effectors. Thus, these methods have to be completely undertaken by their internal control algorithm: this is a result of the embodiment and situatedness prescriptions. In an implementation stage, the notion of "coordi- nation of agents" deals with the organization of the actual processes, or "pieces of software" (structures, objects, .) which represent the agents at machine level. In a serial implementation, agents work in a round robin fashion in such a way that data consistency is preserved. Therefore, no further coordination problems occur. Thus, serial implementations do not to perturb the conceptual definition of coordination of the agents. But in a distributed implementation, new problems arise, due to shared resources and data, synchronization and consistency concerns. Coordination models, in distributed computing, are aimed at providing solutions for these problems: they describe coordination media and tools, external to the agents, so as to deal with the consequences of the spatial distribution of the supporting processes. For example, the coordination medium, such as ports and connections in ECM, is the substratum in which coordinated entities are embedded for what concerns their "coordi- nation dimension". This substratum should not be confused with the agents' environment described at MAS level. It reflects the distributed supporting ar- chitecture. We find here again the notion of orthogo- nality: the coordination medium is orthogonal to the model of the agents and their environment. If we take the separation of concern into account, this means that the coordination introduced at the conceptual level, which has to be implemented in the control algorithm module of the agents, belongs to the computational part, and has to be implemented in the computational language, whereas the coordination of the supporting processes has to be implemented in the coordination language. Nevertheless, the question is now to determine to which extent this separation between computation (including agents' conceptual coordination) and co-ordination (of distributed processes) is possible. In other words, to which extent coordination methods of distributed processes do not interfere with coordination methods of agents as specified at MAS level; and what kind of coordination media and tools may provide coordination means compatible with the agent architecture, their autonomy and locality prescriptions Usual platforms that enable distributed computing do not belong to MAS domain. Distributed implementations of a MAS through existing languages would give rise, if no precautions are taken, to a hybrid system which realizes an improper junction between the two levels of definition of coordination. In each case, the agent processes are coordinated or synchronized at a rate and by means out of the conceptual definition of the MAS, but that the chosen language provides and compels to use in order to manage its processes. Because the processes are not designed with the aim of representing autonomous agents, the resulting system may exhibit characteristics which are not the image of a property of the MAS itself. One of our goals is to understand and prescribe what precautions are to be taken, and to develop a platform that makes distributed implementations of our MAS class an easier task. For this, we start with an implementation of our MAS class with STL. 5.4 Constraints for a Distributed Im- plementation Our very aim is to be able to express our autonomy- based multi-agent model on a distributed architecture in the most natural way which preserve autonomy and identity of the agents. We attempt to use STL in order to distribute our system of gathering agents. The problematics sketched here above is well reflected when we try to distribute the environment itself on several processes (machines). The only purpose of this division of the environment (for instance 4 blocks of (L=2) 2 cells each) is to take advantage of a given distributed architecture. But it clearly necessitates means in addition to coordination mechanisms described at MAS level: a mechanism is needed to cope with agents crossing borders between sub-environments (of course this should be achieved transparently to the user, it should be part of the software platform). Moreover, we will need another type of mechanism in order to cope with data consistency (e.g. updating the number of objects on a cell). These mechanisms should not alter every agen- t's autonomy and behavior: we will have to dismiss any unnecessary dependency. 5.5 Implementation in STL The Environment is a torus grid, in which every cell has four neighbors (four connectivity). Note that using a four connectivity (against an eight connectiv- ity) basically does not change anything except that it slightly alleviates the implementation. Agents comply rigorously with the model previously introduced in Figure 7. They sense the environment through their sensors and act upon their perception at once. To take advantage of distributed systems, the Environment is split into sub-environments, each of which being encapsulated in a blop, as depicted in Figure 9, thus providing an independent functioning between sub-environments (and hence between agents roaming in different sub-environments). Note that blops have to be arranged so as to preserve the topology of the sub-environments they implement. Figure 9: Environment split up among 4 blops. For our implementation new port types have been introduced, namely P2P Nin and P2P Nbi, which are respectively variants of P2Pin and P2P, for which the saturation attribute is set to infinity (see Figure 10). port P2P-Nin - port P2P-Nbi - Figure 10: User-defined port types. Global Structure The meta-blop world is composed of an init process, in charge of the global initialization of the system, and a set of N pre-defined blops (called bx with x ranging from 1 to N), each of which encapsulating and handling a sub-environment. Figure 11 gives a graphical overview of the organization within a blop bx in the meta-blop world. Note that on this figure, only one blop bx is represented so as to avoid to overload the picture. In the case of an application with multiple blops bx, there should be some connections between the init process and all the blops, as well as some connections between north ports of top blops with south ports of bottom blops and east ports of east blops with west ports of west blops have been intentionally dropped The init process has four static ports for every blop to be initialized: three of type P2Pout (init NbAgts, cre Agts and cre SubEnv) and one of type P2P (eot). The r-ole of the init process is threefold: first, to create through its init NbAgts and cre Agts ports the initial agents within every blop; second, to set up through its cre SubEnv port the sub-environment (size, number of objects and position of the objects on the cells) of every blop; third, to collect the result of an exper- iment, to signal the end of an experiment, and to properly shutdown the system through the eot port. Blops bx: Figure 13 (see Appendix) shows how the implementation of the application depicted on Figure looks like in the STL coordination lan- guage, in the case of four blops bx, namely b1, b2, b3 and b4. Figure 12 (see Appendix) presents the declaration and instantiation of all the processes belonging to a blop bx. Two types of processes may be distinguished: processes that are purpose-built for a distributed implementation of the multi-agent application (they enable a distributed implementation), namely initAgent and taxi, and processes that are actually peculiar to the multi-agent application, viz. subEnv and agent processes. Ports of a Blop bx: Each blop has twelve static ports: four P2Pout outflowing direction ports (north o, south o, west o, east o) and four P2Pin inflowing direction ports (north i, south i, west i, east i), which are gateway ports enabling agent migration across blops; three P2Pin ports, namely Agents and b SubEnv used for the creation of the initial agents (actually realized in the initAgent process) and for an appropriate setup of the sub-environment (achieved in the subEnv process); and a P2P port (b eot) used to forward to the init process the result of an experiment and to indicate the end of an experiment. For the time being, the topology between blops is set in a static manner, by creating the ports with appropriate names (see Figure 13 of the Appendix). The four inflowing direction ports of a blop match with a port of its inner process initAgent. The four outflowing direction ports of a blop match with ports of its inner process taxi. initAgent Process, new agt Event: The initA- gent process is responsible for the creation of the agents. It has four static ports: nb Agts of type P2Pin, newArrival of type P2P Nin, location of type P2P and init of type P2Pout. At the outset of the ex- periment, the initAgent process through its nb Agts port will be informed by the init process on the number of agents to be created in the present blop. The initAgent process will then loop on its newArrival port so as to receive the identifiers of the agents to be created. As soon as a value comes to this port, the new agt event (see Figure 12 of the Appendix) is triggered and it will create a new agent process. In the meantime, the initAgent process will draw randomly for every agent's identifier an agent's position. The location port enables the initAgent process to communicate with the subEnv process, so as to have a better control on the position of an agent with regard to other agents' and objects' positions, e.g. to ensure that at the outset no more than one agent can reside on an empty cell. The initAgent process will then write on its init port some values for the agent just created. The latter, through its creation port will read the information that was previously written on the init port of the initAgent process. Values that are transmitted feature for instance the position of the agent and its state. Note that the newArrival port is connected to all inflowing direction ports of the blop within which it resides, thus enabling to deal with migrating agents across blops in the course of an experiment, by the same event mechanisms as described above. The lo- init cre_Agts_b1 cre_SubEnv_b1 Blop bx b_Agents creation req-ans Blackboard b_north_o b_west_o b_south_o b_east_o taxi tSouth tWest con_Agt tNorth tEast agent creation req_ans initAgent init newArrival subEnv init location location nb_Agts eot b_eot b_NbAgts World Figure 11: init process and a single blop bx: solid and dotted lines are introduced just for a purpose of visualization. cation port is very useful at this point, because the initAgent process can already check with the subEnv process whether the position the agent intends to move to is permitted or not. In case it is not, the initAgent process will have to draw randomly a new position in the neighborhood of the position the agent intended to go. agent Process: This process has two static ports (req ans of type BB and creation of type P2Pin) plus to taxi a dynamic P2Pout port. As already stated, this process reads on its creation port some values (its position and its state). All req ans ports of the agents are connected to a Blackboard, through which agents will sense their environment (perception) and act (action) into it, by performing Linda-like in/out operations with appropriate messages. The type of action an agent can take depends on the type of control Algorithm implemented within the agent (see the architecture of an agent on Figure 7). The to taxi port is used to communicate dynamically with the taxi process in case of migration: the position and state of the agent are indeed copied to the taxi pro- cess. The decision of migrating is always taken by the subEnv process. subEnv Process: The subEnv process handles the access to the sub-environment and is in charge of keeping data consistency. It is also responsible for migrating agents, which will cross the border of a sub- environment. It has a static in out port (of type BB) connected to the Blackboard and a static P2Pout port to taxi connected to the taxi process. Once initialized through its init P2Pin port, the subEnv process builds the sub-environment. By performing in/out operations with appropriate tuples, the subEnv process will process the requests of the agents (e.g. number of objects on a given cell, move to next cell) and reply to their requests (e.g. x objects on a given cell, move allowed and registered). When the move of an agent will lead to cross the border (cell located in another blop), the subEnv process will first inform the agent it has to migrate and then inform the taxi process an agent has to be migrated (the direction the agent has to take will be transmitted, so that the taxi process can know which port to write to). The location P2P port is used to communicate to the initAgent process further to its request on the position of an agent with regard to other agents' and objects' positions. The taxi Process: The taxi process is responsible for migrating agents across blops. It has four static direction ports (of type P2Pout), which are connected to the four outflowing direction ports of the blop within which it stands. When this process receives on its static P2Pin port requ the direction towards where an agent has to migrate, it will create a dynamic P2Pin port con Agt in order to establish with the appropriate agent process a communication, by means of which it will collect all the useful information of the agent (intended position plus state). These values will then be written on the port corresponding to the direction to take and will be transferred to the newArrival port of the initAgent process of the concerned blop inducing the dynamic creation of a new agent process in the blop, thus materializing the migration. 6 Discussion 6.1 STL a coordination language As a coordination language for distributed program- ming, ECM along with STL present some similarities with several coordination languages, and particularly with the IWIM model [2] and its instantiation Manifold [3]. However they differ in several important points: ffl One might be inclined to identify blops with IWIM managers (manifolds). This is not the case, because blops are not coordinators that create explicitly interconnections between ports. The establishment of connections is implicit, resulting from a matching mechanism, depending on the types and the states of the ports. This is definitely a different point of view in which communication patterns are not imposed. Further- more, the main characteristics of blops is to encapsulate objects, thus forming a separate name-space for enclosed entities and an encapsulation mechanism for events. Nested blops are a powerful mechanism to structure private name-spaces, offering an explicit hierarchical model. ffl ECM generalizes connection types: either stream, blackboard or group. This adds powerful means to express coordination with tuple-space models and does not restrict to channels. Refined semantics can be defined in virtue of port characteristics (features). ffl In ECM, events are not signals broadcast in the environment, but routines belonging to blops. They are attached to ports with conditions on their state that determine when events are launched. Events can create new blops and pro- cesses, and attach events to ports. Their action area is limited to a blop. ffl Interconnections evolve through configuration changes of the set of ports within a blop, induced by events and also by the processes them- selves. In fact, the latter can create new processes and new ports, thus yielding to communication topology changes. ECM and STL present similarities with several other coordination models and languages like Linda [9], Darwin [24] or ConCoord [18]. We mention few other specific characteristics of our work. Like several further developments of the Linda model (for instance Objective Linda [19]), ECM uses a hierarchical multiple coordination space model, in contrast to the single flat tuple space of the original Linda. Processes get started through an event in a blop, or automatically upon initialization of a blop, or through a creation operation by another process; Linda uses one single mechanism: eval(). Processes do not execute in a medium which is used to transfer data. In order to communicate, they do not have references to other processes or to ports belonging to other entities; they communicate anonymously through their ports. 6.2 The agent language STL++, a new instantiation of the ECM model Besides the advantage of a better overview of coordination duties, it however turned out that the separation of code (as in STL) can not always be main- tained. Although the black box process model of ECM is a good attempt to separate coordination and computation code, dynamic properties proved to be difficult to express in a separate language. This is for example reflected in STL by the primitives which must be used in the computation language in order to use dynamic coordination facilities of STL. Dynamic properties can not be separated totally from the actual program code. Furthermore, a duplication of code for processes may introduce difficulties to manage code for a distributed application. These observations lead us to the development of another instantiation of ECM, namely the coordination language STL++. This new language binding implements ECM by enriching a given object oriented language (C++ in our case) with coordination primitives, offering high dynamical properties. An STL++ application is then a set of classes inheriting from the basic classes of the STL++ library. STL++ aims at giving basic constructs for the implementation of generic multi-agent platforms, thus being an agent language [32]. A thorough description of STL++ can be found in [29]. 7 Conclusion In this paper, we presented the ECM coordination model and STL, its language binding. We built a first STL-based prototype on top of the existing Pt- pvm platform [22]. An implementation of a classical collective robotics simulation illustrated the power of and demonstrated its appropriateness for coordinating a class of autonomous agents, whose most critical constraint is the preservation of autonomy by dismissing coordination mechanisms exclusively embedded for purpose of implementation (unnecessary dependencies). As far as the development of a platform for multi-agent programming is concerned, STL can be seen as a first starting point. STL already includes mechanisms which are appropriate for multi-agent pro- gramming, among which are: (1) the absence of a central coordinator process, which does not relate to any type of entity in the multi-agent system; (2) the notion of ports avoiding any additional coordinator process; and (3) in despite of (2) the notion of blop hierarchy which in our case allows us to represent the encapsulation of the environment and the agents. The STL coordination model is still to be extended in order to encompass as many generic coordination patterns as possible, yielding in STL skeletons at disposal for general purpose implementations. Future works will consist in: (1) improving the model, such as introducing new user-defined attributes for ports, dynamic ports for blops, data typing for port types, refining sub-typing of ports, supporting multiple names for ports, and (2) developing a graphical user interface to facilitate the specification of the co-ordination part of a distributed application. There are two major outcomes to this work. First, as autonomous agents' systems are aimed at addressing problems which are naturally distributed, our coordination platform provides a user the possibility to have an actual distributed implementation and therefore to benefit from the numerous advantages of distributed systems, so that this work is a step forward in the Autonomous Agents commu- nity. Secondly, as the generic patterns of coordination for autonomy-based multi-agent implementations are embedded within the platform, a user can quite easily develop new applications (e.g. by changing the type of autonomy of the agents, the type of insofar they comply with the generic model. Acknowledgements We are grateful to Andr'e Horstmann and Christian Wettstein for their valuable work, which consisted in realizing parts of the STL platform. --R ACTORS: A Model of Concurrent Computation in Distributed Systems. The IWIM Model for Coordination of Concurrent Activities. An Overview of Manifold and its Implementation. Programming by Multiset Transformation. The ToolBus Coordination Architecture. Extending Objects with Rules Intelligence without Reason. Adaptive Parallelism and Pi- ranha Linda in Con- text Coordination Languages and Their Significance. Bauhaus Linda. Coop'eration implicite et Autonomous Agents: from Concepts to Implementation. An Architecture to Co-ordinate Distributed Applications on the Web Performance of Autonomy-based Systems: Tuning Emergent Cooperation CoLa: A Coordination Language for Massive Parallelism. A Software Environment for Concurrent Coordinated Programming. Objective Linda: A Coordination Model for Object-Oriented Parallel Pro- gramming Coordination Requirements for Open Distributed Systems. and Pt-PVM: Concepts and Tools for Coordination of Multi-threaded Appli- cations Artificial Life. Structuring parallel and distributed programs. The Interdisciplinary Study of Coordination. The Design of Intelligent Agents: A Layered Approach. Coordination Models and Languages. A Set of Tuple Space primitives for Distributed Co- ordination Programming Languages for Parallel Processing. Agent The- ories Adaptive Behavior in autonomous agents. --TR

autonomous agents;distributed systems;coordination;collective robotics

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Supporting Scenario-Based Requirements Engineering.

AbstractScenarios have been advocated as a means of improving requirements engineering yet few methods or tools exist to support scenario-based RE. The paper reports a method and software assistant tool for scenario-based RE that integrates with use case approaches to object-oriented development. The method and operation of the tool are illustrated with a financial system case study. Scenarios are used to represent paths of possible behavior through a use case, and these are investigated to elaborate requirements. The method commences by acquisition and modeling of a use case. The use case is then compared with a library of abstract models that represent different application classes. Each model is associated with a set of generic requirements for its class, hence, by identifying the class(es) to which the use case belongs, generic requirements can be reused. Scenario paths are automatically generated from use cases, then exception types are applied to normal event sequences to suggest possible abnormal events resulting from human error. Generic requirements are also attached to exceptions to suggest possible ways of dealing with human error and other types of system failure. Scenarios are validated by rule-based frames which detect problematic event patterns. The tool suggests appropriate generic requirements to deal with the problems encountered. The paper concludes with a review of related work and a discussion of the prospects for scenario-based RE methods and tools.

Introduction Several interpretations of scenarios have been proposed ranging from examples of behaviour drawn from use cases [29], descriptions of system usage to help understand socio-technical systems [30], and experience based narratives for requirements elicitation and validation [39], [50]. Scenarios have been advocated as an effective means of This research has been funded by the European Commission ESPRIT 21903 'CREWS' (Co-operative Requirements Engineering With Scenarios) long-term research project. communicating between users and stakeholders and anchoring requirements analysis in real world experience [15]. Unfortunately scenarios are extremely labour-intensive to capture and document [20], [44]; furthermore, few concrete recommendations exist about how scenario-based requirements engineering (RE) should be practised, and even less tool support is available. Scenarios often describe information at the instance or example level. This raises the question of how instance level information can be generalised into the models and specifications that are used in software engineering. Scenarios may be used to validate requirements, as 'test data' collected from the observable practice, against which the operation of a new system can be checked [40]. Alternatively, scenarios may be seen as pathways through a specification of system usage, and represented as animations and simulations of the new system [14]. This enables validation by inspection of the behaviour of the future system. This paper describes a method and tool support for scenario based requirements engineering that uses scenarios in the latter sense. In industrial practice scenarios have been used as generic situations that can prompt reuse of design patterns [8], [48]. Reuse of knowledge during requirements engineering could potentially bring considerable benefits to developer productivity. Requirements reuse has been demonstrated in a domain specific context [31], however, we wish to extend this across domains following our earlier work on analogies between software engineering problems [51]. The paper is organised in four sections. First, previous research is reviewed, then in section 3 a method and software assistant tool for scenario based RE, CREWS-SAVRE (Scenarios for Acquisition and Validation of Requirements) is described and illustrated with financial dealing system case study. Finally, we discuss related work and future prospects for scenario based RE. 2. Previous Work Few methods advise on how to use scenarios in the process of requirements analysis and validation. One of the exceptions is the Inquiry Cycle of Potts [39] which uses scenario scripts to identify obstacles or problems in a goal-oriented requirements analysis. Unfortunately, the Inquiry Cycle does not give detailed advice about how problems may be discovered in scenarios; furthermore, it leaves open to human judgement how system requirements are determined. This paper builds on the concepts of the Inquiry Cycle with the aim of providing more detailed advice about how scenarios can be used in requirements validation. In our previous work we proposed a scenario based requirements analysis method (SCRAM) that recommended a combination of concept demonstrators, scenarios and design rationale [50]. SCRAM employed scenarios scripts in a walkthrough method that validated design options for 'key points' in the script. Alternative designs were documented in design rationale and explained to users by demonstration of early prototypes. SCRAM proved useful for facilitating requirements elaboration once an early prototype was in place [48]; however, it gave only outline guidance for a scenario-based analysis. Scenarios can be created as projections of future system usage, thereby helping to identify requirements; but this raises the question of how many scenarios are necessary to ensure sufficient requirements capture. In safety critical systems accurate foresight is a pressing problem, so taxonomies of events [24] and theories of human error [37], [42], have been used to investigate scenarios of future system use. In our previous research [49], we have extended the taxonomic approach to RE for safety critical systems but this has not been generalised to other applications. Scenarios have been adopted in object oriented methods [29], [21], [10] as projected visions of interaction with a designed system. In this context, scenarios are defined as paths of interaction which may occur in a use case; however, object oriented methods do not make explicit reference to requirements per se. Instead, requirements are implicit within models such as use cases and class hierarchies. Failure to make requirements explicit can lead to disputes and errors in validation [17], [38]. Several requirements management tools have evolved to address the need for requirements tracability (e.g. Doors, Requisite Pro). Even though such tools could be integrated with systems that support object oriented methods (e.g. Rational Rose) at the syntactic level, this would be inadequate because the semantic relationships between requirements and object oriented models needs to be established. One motivation for our work is to establish such a bridge and develop a sound means of integrating RE and OO methods and tools. Discovering requirements as dependencies between the system and its environment has been researched by Jackson [26] who pointed out that domains impose obligations on a required system. Jackson proposed generic models, called problem frames, of system-environment dependencies but events arising from human error and obligations of systems to users were not explicitly analysed. Modelling relationships and dependencies between people and systems has been investigated by Yu and Mylopoulos [54] and Chung [7]. Their i* framework of enterprise models facilitates investigation of relationships between requirements goals, agents and tasks; however, scenarios were not used explicitly in this approach. Methods are required to unlock the potential of scenario based RE; furthermore, the relationship between investigation based on examples and models on one hand and the systems and requirements imposed by the environment on the other, needs to be understood more clearly. One problem with scenarios is that they are instances, i.e. specific examples of behaviour which means that reusing scenario based knowledge is difficult. A link to reusable designs might be provided if scenarios could be generalised and then linked to object-oriented analysis and design patterns [18]. Although authors of pattern libraries do describe contexts of use for their patterns [8], [19], they do not provide guidance about the extent of such contexts. Requirements reuse has been demonstrated by Lam et al. [30]; although, the scope of reuse was limited to one application domain of jet engine control systems. Many problems share a common abstraction [51], and this raises the possibility that if common abstractions in a new application domain could be discovered early in the RE process, then it may be possible to reuse generic requirements and link them to reusable designs. This could provide a conduit for reusing the wealth of software engineering knowledge that resides in reusable component libraries, as well as linking requirements to object-oriented solution patterns (e.g. [8], [19]). Building a bridge from requirements engineering to reusable designs is another motivation for the research reported in this paper. Next we turn to a description of the method and tool support. 3. Method and Tool support for Scenario-based RE The method of scenario-based RE is intended to be integrated with object oriented development (e.g. OOSE - [29]), hence use cases are employed to model the system functionality and behaviour. A separate requirements specification document is maintained to make requirements explicit and to capture the diversity of different types of requirements, many of which can not be located in use cases. Use cases and the requirements specification are developed iteratively as analysis progresses. We define each scenario as "one sequence of events that is one possible pathway through a use case". Hence many scenarios may be specified for one use case and each scenario represents an instance or example of events that could happen. Each scenario may describe both normal and abnormal behaviour. The method illustrated in Figure 1 uses scenarios to refine and validate system requirements. The stages of the method are as follows: 1. Elicit and document use case In this stage use cases are elicited directly from users as histories of real world system usage or are created as visions of future system usage. The use case model is validated for correctness with respect to its syntax and semantics. The paper does not report this stage is in detail. 2. Analyse generic problems and requirements A library of reusable, generic requirements attached to models of application classes is provided. A browsing tool matches the use case and input facts acquired from the designer to the appropriate generic application classes and then suggests high level generic requirements attached to the classes as design rationale 'trade-offs'. Generic requirements are proposed at two levels: first, general routines for handling different event patterns, and secondly, requirements associated with application classes held in the repository. The former provide requirements that develop filters and validation processes as specified in event based software engineering methods (e.g. Jackson System Development [25]); while the latter provide more targeted design advice; for instance, transaction handling requirements for loans/hiring applications. 3. Generate scenarios This step generates scenarios by walking through each possible event sequence in the use case, applying heuristics which suggest possible exceptions and errors that may occur at each step. This analysis helps the analyst elaborate pathways through the use case in two passes; first for normal behaviour and secondly for abnormal behaviour. Each pathway becomes a scenario. Scenario generation is supported by a tool which automatically identifies all possible pathways through the use case and requests the user to select the more probable error pathways. 4. Validate system requirements using scenarios Generation is followed by tool-assisted validation which detects event patterns in scenarios and presents checklists of generic requirements that are appropriate for particular normal and abnormal event patterns. In this manner requirements are refined by an interactive dialogue between the software engineer and the tool. The outcome is a set of formatted use cases scenarios and requirements specifications which have been elaborated with reusable requirements. Although the method appears to follows a linear sequence, in practice the stages are interleaved and iterative. Figure 1 Method Stages for Scenario-based Requirements Engineering. 3.1. Schema for scenario based requirements modelling In this section we describe the schema of scenario based knowledge shown in Figure 2. A use case is a collection of actions with rules that govern how actions are linked together, drawn from Allen's temporal semantics [2] and the calculus of the ALBERT II specification language [14]. An action is the central concept in both scenarios and use cases. The use case specifies a network of actions linked to the attainment of a goal which describes the purpose of the use case. Use cases are refined into lower levels of detail, so a complete analysis produces a layered collection of use cases. Use cases have user-defined properties that indicate the type of activity carried out, e.g. decision, transaction, control, etc. Each action can be sub typed as either cognitive (e.g. the buyer agrees the deal price on A'.h#vt #@yvpv#"+r ph+r !#6hy'+r Brr.vp Whyvqh#v'A.hr+ "#Brr.h#r #Whyvqh#r Key: offer), physical (e.g. the dealer returns the telephone receiver), system-driven (e.g. the dealing systems stores information about the current deal) or communicative (e.g. the buyer requests a quote from the dealer). Each action has one start event and one end event. Actions are linked through 8 types of action-link rules which are described later in more detail. state transition state object action use case agent contains starts/ ends results in from /to has event uses involves scenario is sequence of goal achieves model primitives models generated from structure object composed property has 1.m ipates in Figure 2 Meta-schema for modelling use cases and scenario based knowledge. Each action involves one or more agents. Each agent can be either human (e.g. a dealer), machine (e.g. a dealing-system), composite (e.g. a dealing-room) or an unspecified type. Agents have user-defined properties that describe their knowledge and competencies. Action - agent links are specified by an involvement relation that can be subtyped as {performs, starts, ends, controls, responsible}. Each action uses nil, one or many object instances denoted by a use relation, subtyped as {accesses, reads, operates}. Each action can also result in state transitions which change the state of objects. States are aggregations of attribute values which characterise an object at a given time in both quantitative and qualitative terms. Structure objects are persistent real world objects that have spatial properties and model physical and logical units of organisation. Structure objects are important components in reusable generic domain models (called object system models and described in section 3.5), but also allow use case models to be extended to describe the location of objects, agents and their activity. Introduction to the case study The case study is based on a security dealing system at a major bank in London. Securities dealing systems buy and sell bonds and gilt-edged stock for clients of the bank. The bank's dealers also buy and sell securities on their own account to make a profit. Deals may be initiated either by a client requesting transactions, or by a counterparty (another bank or stockbroker who acts as a buyer) requesting a quotation, or by the dealer offering to buy or sell securities. Quotes are requested over the telephone, while offers are posted over electronic wire services to which other dealers and stockbrokers subscribe. In this case study we focus on deals initiated by a counterparty from the viewpoint of the dealer. The main activities in this use case include agreement between the dealer and buyer on the deal price and number of stock, entering and recording deal information in the computer system, and checking deal information entered by the dealer to validate the deal. The case study contains several use cases; however, as space precludes a complete description, we will take one use case ' prepare-quote' as an example, as illustrated in figure 3. The sequence is initiated by a request event from the counterparty agent. The dealer responds by providing a quotation which the counterparty assesses. If it is suitable the counterparty agrees to the quotation and the deal is completed. Inbound events to the system are the deal which has to be recorded and updated, while outbound events are displays of market information and the recorded deal. System actions are added to model the first vision of how the dealing support system will work. The dealer agent is linked to the dealing room structure object that describes his location of work. The dealer carries out the "prepare quote" use case which is composed of several actions and involves the "trade" object. Figure 3 Upper level use case illustrated as an Agent-interaction showing tasks, agents and actions. Counter-Party Dealing System Dealer Head Dealer 3.2. Tool support The method is supported by Version 2.1 of the CREWS-SAVRE tool that has been developed on a Windows-NT platform using Microsoft Visual C++ and Access, thus making it compatible for loose integration with leading commercial requirements management and computer-aided software engineering software tools. It supports 6 main functions which correspond to the architecture components shown in Figure 4. 1: incremental specification of use cases and high-level system requirements (the domain/use case modeller supports method stage 1) 2: automatic generation of scenarios from a use case (scenario generator supports stage 3: manual description of use cases and scenarios from historical data of previous system use, as an alternative to tool-based automatic scenario generation (use case/scenario authoring component supports stage 1); 4: presentation of scenarios, supporting user-led walkthrough and validation of system requirements (scenario presenter supports stage 4); 5: semi-automatic validation of incomplete and incorrect system requirements using commonly occurring scenario event patterns (requirements validator supports stage 4). CREWS-SAVRE is loosely coupled with RequisitePro's requirements database so it can make inferences about the content, type and structure of the requirements. Another component, which guides natural language authoring of use case specifications, is currently under development. The component uses libraries of sentence case patterns (e.g. [14]) to parse natural language input into semantic networks prior to transforming the action description into a CREWS-SAVRE use case specification. Space precludes further description of the component in this paper. The CREWS-SAVRE tool permits the user to develop use cases which describe projected or historical system usage. It then uses an algorithm to generate a set of scenarios from a use case. Each scenario describes a sequence of normal or abnormal events specified in the original use case. The tool uses a set of validation frames to detect event patterns in scenarios, thereby providing semi-automatic critiquing with suggestions for requirements implied by the scenarios. domain /use case modeller scenario generator scenario presenter requirements validater use case author tool CREWS-SAVRE tool REQUISITEPRO REQUIREMENTS MANAGEMENT scenario author tool natural language descriptions use case facts scenario facts generated scenarios generated scenarios scenarios validated requirements user/ domain expert use case environment modeller/ validater system-env models user/ software engineer Figure 4 Overview of the CREWS-SAVRE tool architecture 3.3 Use Case Specification The user specifies a use case using CREWS-SAVRE's domain and use case modeller components. First, for a domain, the software engineer specifies all actions in the domain, defines agents and objects linked to these actions, assigns types {e.g. communicative, physical, etc.} to each action, and specifies permissible action sequences using action-link rules. From this initial domain model, the user can then choose the subset of domain actions which form the normal course of a use case. This enables a user to specify more than one use case for a domain, and hence generate more scenarios for that domain. As a result, a use case acts as a container of domain information relevant to the current scenario analysis. As scenarios are used to validate different parts of the domain's future system, different use cases containing different action descriptions are specified. Consequently, the domain model enables simple reuse of action descriptions across use cases, because two or more use cases can include the same action. Eight types of action-link rule are available in CREWS-SAVRE: strict sequence part sequence inclusion includes B): (ev(startA)<ev(startB)) AND ((ev(endA)>(ev(endB)); concurrent no rules about ordering of events; alternative parallel and equal equal-start starts-with B): (ev(startA)=ev(startB)) AND ((ev(endA)not=(ev(endB)); equal-end where event ev(X) represents the point in time at which the event occurs. Actions are defined by ev(startA)<ev(endA) and ev(startB)<ev(endB), that is all actions have a duration and the start event of an action must occur before the end event for the same action. These link rules types build on basic research in temporal semantics [2] and the formal temporal semantics and calculus from a real-time temporal logic called CORE which underpins the ALBERT II specification language [23]. The current version of CREWS-SAVRE does not provide all of Allen's [2] 13 action-link types. Rather, it provides a set of useful and usable semantics based on practical reports of use case analysis (e.g. [1], [29]). For the dealing system domain, two actions are Buyer requests deal from the dealer and Dealer retrieves price information from the dealer-system. Selection of the action-link rule MEANWHILE indicates that, in general, the dealer begins to retrieve price information only after the buyer begins to request the deal. A third action, this time describing system behaviour- Dealer-system displays the price information to the dealer- can be linked through the THEN rule to specify a strict sequence between the two actions, in that the first action must end before the second action starts. Part of the dealing system domain model is shown in Figure 5. It shows a subset of the current domain actions (shown as (A) in Figure 5), the specification of attributes of a new action (B), current action-link rules (C) and agent types (D). Other parts of the domain modeller are outside of the scope of this description. One or more use cases are specified for a domain such as the dealing system. Each use case is linked to one high-level requirement statement, and each system action to one or more system requirements. Each use case specification is in 4 parts. The first part specifies the identifiers of all actions in the use case. The second part specifies action-link rules linking these actions. The third part contains the object-mapping rules needed to handle use of synonyms in action descriptions. The fourth part specifies exception-types linked to the use case, as described in section 3.6. Figure 5 Domain Modeller screen dump. 3.4 Checking the use case specification Validation rules check the integrity of the relationships specified in section 3.1, represented in tuple format < model component1, relationship-type, model component2>. The tool checks the integrity of use case models to ensure they conform to the schema, but in addition, the tool is parameterised so it can check use cases in a variety of ways. This enables user defined validation of properties not defined in the schema. Validation checks are carried out by clusters of rules, called validation-frames which are composed of two parts. First a situation that detects a structural pattern (i.e. a combination of components connected by a relationship) in the use case. The second part contains requirements that should be present in the part of the model detected by the situation. Frames are used for validating the consistency of use case models against the schema, and for detecting potential problems in use cases as detailed in section 3.5. In the former case the frames detect inconsistencies in a use case, the latter case frames detect event patterns in scenarios and suggest appropriate, generic requirements. Two examples of consistency checking frames are as follows: (i) Checks for agents connected by a specific relationship type. validation-frame {detects dyadic component relationships} situation: model(component(x), relationship (y), component (z)); schema requirements: (component type (i) component (x), mandatory) ; (component type (j) component (y), mandatory); end-validation-frame Example: The user decides to check that all actions are controlled by at least one agent. First the tool finds all nodes with a component type = agent, then it finds all the nodes connected to that agent which are actions and finally whether the involvement relationship that connects the two nodes is of the correct As the tool is configurable the input parameters, (i) and (j) can be any schema primitives, and the relationship (k) is defined in the schema. The tool detects untyped or incorrect relationships using this 'structure matching' algorithm. In the 'prepare quotes' use case (see figure 3 ), this validation check should detect that the dealer controls the give-quote action. (ii) Validates whether two components participating in a relationship have specific properties. validation-frame {detects components in a relationship have the correct properties} situation: model(component(x), property(w), relationship (z), component (y), property (v)); schema requirements: (component type (i) component (x), mandatory); (component type (j) component (y), mandatory); (component (x) property (w), mandatory); (component (y) property (v), mandatory); end-validation-frame Example: The user wishes to validate that all agents that are linked to a use case with a property have an property. The properties (v,w) to be tested are entered in a dialogue, so any combination may be tested. In this case the tool searches for nodes with a type = agent, and then tests all the relationships connected to the agent node. If any of these relationships is connected to a node type = use case, then the tool reads the property list of the use case and the agent. If the agent does not have an 'authority' property and the use case has a 'decision' property then warning is issued to the user. As properties are not sub types, this is a lexical search of the property list. This frame will test that the dealer has authority for the use case 'evaluate choice' - which has a decision property. The system is configurable so any combinations of type can be checked for any set of nodes and arcs which conform with the schema described in figure 2. Our motivation is to create requirements advice that evolves with increasing knowledge of the domain, so the user can impose constraints beyond those specified in the schema, and use to tool to validate that the use case conforms to those constraints. 3.5 Using Application classes to identify Generic requirements This stage takes the first cut use case(s) and maps them to their corresponding abstract application classes. This essentially associates use cases describing a new system with related systems that share the same abstraction. A library of application classes, termed Object system models (OSMs) has been developed and validated in our previous work on computational models of analogical structure matching [34]. Object system models (OSMs) are organised in 11 families and describe a wide variety of applications classes. The families that map to the case study application are object supply (inventory control), accounting object transfer (financial systems), object logistics (messaging), object sensing (for monitoring applications), and object-agent control (command and control systems). Each OSM is composed of a set of co-operating objects, agents and actions that achieve a goal, so they facilitate reuse at the system level rather than as isolated generic classes. Taking object supply (see Appendix A) as an example, this OSM models the general problem of handling the transaction between buyers and suppliers. In this case study, this matches to the purchase of securities which are supplied by the bank to the counter party who acts as the buyer. The supplier giving a price maps to the dealer preparing a quotation. Essentially OSMs are patterns for requirements specification rather than design solutions as proposed by [19]. Each OSM represents a transaction or a co-operating set of objects and are modelled as class diagrams, while the agent's behaviour is represented in a use case annotated with high level generic requirements, expressed in design rationale diagrams. A separate set of generic use case models are provided for the functional or task-related aspects of applications. Generic use cases cause state changes to objects or agents; for instance, a diagnostic use case changes the state of a malfunctioning agent (e.g. human patient) from having an unknown cause to being understood (symptom diagnosed). Generic use cases are also organised in class hierarchies and are specialised into specific use cases that are associated with applications. Currently seven families of generic use cases have been described; diagnosis, information searching, reservation, matching- allocation, scheduling, planning, analysis and modelling. An example of a generic use case is information searching which is composed of subgoals for: articulating a need, formulating queries, evaluating results, revising queries as necessary. In the dealing domain this maps to searching for background information on companies in various databases, evaluating company profit and loss figures and press releases, then refining queries to narrow the search to specific companies of interest. For a longer description of the generic use cases and their associated design rationale see [53]. Specific use cases in a new application may be associated with object system models either by browsing the OSM/use case hierarchy and selecting appropriate models, or by applying the identification heuristics - see Appendix A, or by using a semi-automated matching tool that retrieves appropriate models given a small set of facts describing the new application [51]. These heuristics point towards OSM models associated with the application; however, identification of appropriate abstractions is complex and a complete description is beyond the scope of this paper. In this stage, mapping between use case components and their corresponding abstractions in the OSMs are identified so that generic requirements, attached to the OSMs, can be applied to the new application. Unfortunately, the mapping of problems to solutions is rarely one to one, so trade -offs have to be considered to evaluate the merits of different solutions. Design rationale [11] provides a representation for considering alternative designs that may be applied to the requirements problems raised by each OSM. Non-functional requirements are presented as criteria by which trade offs may be judged. The software engineer judges which generic requirements should be recruited to the requirements specification and may adds further actions to the use case thereby elaborating the specification. Case study The security trading system involves five OSMs; Object Supply that models securities trading; Account Object Transfer which models the settlement part of the system (payment for securities which have been purchased) and Object messaging to describe the communication between the dealer and counterparties. Other OSMs model sub systems that support trading, such as Object Sensing that detects changes in security prices, markets and the dealer's position, and Agent Control which describes the relationship between the head dealer and the other dealers. In addition to the OSMs, the dealing system contains generic use cases (evaluate purchase and plan strategy) that describe the dealer's decision making and reasoning. These map the domain specific use cases of evaluating a deal that has been proposed and for planning a trading strategy. A further specific use case 'prepare quote' is mapped to the generic use case (price item) associated with the Object Supply OSM. A generic model of the security trading system, expressed as an aggregation of OSMs, is given in figure 6. The settlement part of the system (Accounting object transfer OSM) has been omitted. The OSM objects have been instantiated as dealing system components. Clusters of generic requirements, represented as design rationale are associated with appropriate OSM components. bank security stock client transfer owned-by move request supplier supply dealer banking world wire service signal exist-in quotes change prices head dealer strategies banking world transfer move-in trade move source destination banks quotes make deal, prep quote eval trade plan strategy counter party held-by deal security transfer exist-in price agree eval trade customer request instructions locate Generic requirements clusters (GRs in Design rationales) 1. Checking customer preferences (limits on dealing) 2. Calculating prices (prepare quotations) 3. Evaluating choices (deals) 4. Sampling changes to object properties (stock prices) 5. Message transmission protocols (deal notification) 6. Message encryption (deal security) 7. Communicating commands (head dealer strategy) 8. Reporting compliance (strategy obeyed/followed) 9. Calculating replenishment257 81 Object messaging OSM sub class of Object Logistics Object sensing OSM sub class Object properties Object supply and Agent control OSMs Figure 6 Aggregation of OSMs that match to the security dealing system, represented in object oriented analysis notation [9]. Figure 7 Design rationale for high level generic requirements from clusters 2 and 3 in Figure 6 attached to 'evaluate choice'and 'prepare quote'use cases. The instantiated requirement derived from the generic version is given in brackets. The functional requirements that could be applied to support two use cases 'evaluate deal' and 'prepare quote'. For evaluate deal the rationale is taken from the generic use case 'evaluate choice', which is a sub-class in the matching/allocation family. This proposes three options: to assess the purchase against a set of reference levels, to prioritise several Speed of Operation 9r+vtSh#v'hyrvtD7DTI'#h#v' D++"r Q'+v#v'+ 6.t"r#+ (Generic requirements) Evaluate Choices (deals) Check options Matrix trade-off Constraint-based matching Rapid response Accuracy Sophisticated Choice Complexity/Cost [3] b dC'+r'sR"hyv#'b!!d b"d'#.hqr#'ss#w"+#u.r+u'yqpurpx (design alternatives) (justifications and non-functional Calculate Prices (Prepare Quote) Calculate from volume & value (goods position) Weighted Matrix Calculation Use price file (Display Baseline Quotes) Range of Prices User Choice Positive justification of a position by an argument Negative argument against a position purchase options by a simple House of Quality style matrix [22], and finally to use a sophisticated multi-criteria decision making algorithm. Hypertext links from the rationale point to reusable algorithms for each purpose. The options for 'prepare quotes' are to automate quotation with simple calculations based on the dealer's position, the desirability of the stock and the market movement; or to choose a weighted matrix technique for quoting according to the volume requested and the dealer's position, or to leave quotation as a manual action with a simple display of the bank's baseline quotations. Since the first two options may be too time consuming for the dealers, the third was chosen. For evaluate deal, the simple calculation is taken as optional facility, leaving the dealer in control. Two high level generic requirements are added to the requirements specification and actions to the use case which elaborates the system functionality. 3.6 Scenario Generation This stage generates one or more scenarios from a use case specification. Each scenario is treated as one specific ordering of events. The ordering is dependent on the timings of start- and end- events for each action and the link rules specified in the originating use case. Entering timings is optional, so in the absence of timed events, the algorithm uses the ordering inherent in the link rules. More than one scenario can be generated from a single use case if the action-link rules are not all a strict sequence (i.e. A then B). The space of possible event sequences is large, even within a relatively simple use case. The scenario generation algorithm reduces this space by first determining the legal sequences of actions. The space of permissible event sequences is reduced through application of action-link and user-constraints. The user can enter constraints that specify which agents and actions should be considered in the scenario generation process, thus restricting the permissible event sequences to sequences (es) to include an event (ev) that starts the action (A) and involves a predefined agent (ag) and which has at least a given probability UC: (ev(startA) in es) AND (ev starts For example, each generated scenario must include the event that starts action 20, it must involve the agent "dealer" and action 20 must have at least a 10% likelihood of occurrence according to probabilities calculated from information in the use case specification. The 'prepare quote' use case definition leads to the 3 possible scenarios shown in figure 8. The difference between each is the timing of event E40 which ends action 40, and whether action 40 or action 45 occurs. This depiction ignores the application of the constraint on the likelihood of an event sequence occurring. Scenario-1 and -2 differ in the timing of event E40 (end of request for price information from the dealer-system) while scenario-3 describes a different event sequence when the dealer is unable to offer a quote for the deal. event event (buyer requests quote) event S50 event E50 (dealer- system shows price) Time THEN event event E40 event event event event E40 event S50 event E50 event event event event E40 event S50 event E50 PART OF USE CASE GENERATES event event E20 (dealer picks up THEN event event E20 event event E20 event E40 (dealer retrieves price event S60 event event E45 (dealer refuses deal) THEN event event event event E45 event event E20 Figure 8 Diagram illustrating 3 normal scenario paths generated from a use case fragment The generation mechanism is in two stages. First it generates each permissible normal course scenario from actions and link rules in the use case, then it identifies alternative paths for each normal sequence from exception types, as summarised in Table I. They are divided into two groups. First abnormal events drawn from Hollnagel's [24] event 'phenotypes' classification and secondly, information abnormalities that refer to the message contents(e.g. the information is incorrect, out-of-date etc.) that follow validation concepts proposed by Jackson [27]. Each exception is associated with one or more generic requirements that propose high level solutions to the problem. The exception types are presented as "what-if" questions so the software engineer can choose the more probable and appropriate alternative path at each action step in the scenario. Table I Summary of Exception types for events originating in the system and generic requirements to deal with abnormal patterns. Exception Generic Requirement event does not happen - omitted time-out, request resend, set default event happens twice (not iteration) discard extra event, diagnose duplicate event happens in wrong order buffer and process, too early - halt and wait, too late - send reminder, check task event not expected validate vs. event set, discard invalid event information - incorrect type request resend, prompt correct type incorrect information values check vs. type, request resend, prompt with diagnosis information too late (out of date) check data integrity, date/time check, use default information too detailed apply filters, post process to sort/group information too general request detail, add detail from alternative source A set of rules constrain the generation of alternative courses in a scenario using action and agent types. These rules are part of an extensible set which can be augmented using the agent types and influencing factors described below. Two example rules are: IF ((ev1 starts ac1) OR (ev1 ends ac1)) AND ac1(type=cognitive) THEN (ex(type=human) applies-to ev1). which ensures that only human exception types and influencing factors (see next section) are applied to a cognitive action for event (ev1), action (ac1) and exception type (ex), and IF ((ev1 starts ac1) OR (ev1 ends ac1)) AND ac1(type=communicative) AND (ac1 involves ag1) AND ag1(type=machine) AND (ac1 involves ag2) AND ag2(type=machine) THEN (ex(type=machine-machine-communication) applies-to ev1). which ensures that only machine-machine communication failures are applied to communication actions where both agents are of type 'machine' for event (ev1), action (ac1), agents (ag1 and ag2) and exception type (ex). The rules identify particular types of failure that may occur in different agent-type combinations so that generic requirements can be proposed to remedy such problems. For instance, in the first rule, human cognitive errors that apply to action1 can be counteracted by improved training or aid memoir facilities (e.g. checklists, help) in the system. In the second rule which detects network communication errors, generic requirements are suggested for fault tolerant designs sand back-up communications. The algorithm generates a set of possible alternative paths according to the agent-action combination. The use case modeller allows the user to select the more probable abnormal pathways according to their knowledge of the domain. To help the software engineer anticipate when exceptions may occur and assign probabilities to abnormal events, a set of influencing factors are proposed. These describe the necessary preconditions for an event exception to happen and are sub divided into 5 groups according to the agents involved: human agents: Influencing factors that give rise to user errors and exceptions are derived from cognitive science research on human error [42], Norman's model of slips [37] and Rasmussen's three levels of human-task mismatches [41]. However, as human error cannot be adequately described by only cognitive factors; we have included other performance affecting properties such as motivation, sickness, fatigue, and age; based on our previous research on safety critical systems [49]. machine agents: failures caused by hardware and software, e.g. power supply problems, software crashes, etc. human-machine interaction: poor user interface design can lead to exceptions in input/output operations. This group draws on taxonomies of interaction failures from human-computer interaction [46] and consequences of poor user interface design (e.g. human-human communication: scenarios often involve more than one human agent. Communication breakdowns between people have important consequences. Exceptions have been derived from theories from computer-supported collaborative work [46]. Examples include communication breakdowns and misunderstandings; machine-machine communication: scenarios often involve machine agents, and exceptions specific to their communication can also give rise to alternative paths. The interaction between influencing factors that give rise to human error is described in figure 9. Four outer groups of factors (working conditions, management, task/domain and personnel qualities) effect four inner factors (fatigue, stress, workload and motivations). These in turn effect the probability of human error which is manifest as an event exception of type <human-machine action or human action>. Human error can be caused by environment factors and qualities of the design, so two further outer groups are added. Personnel/user qualities are causal influences on human operational error, whereas the system properties can either be treated as causal explanations for errors or viewed as generic user interface requirements to prevent such errors. Requirements to deal with problem posed by influencing factors are derived from several sources in the literature, e.g. for task design and training [3], workplace ergonomics [45] and for Human Computer Interface design [46], [47] and standards (e.g. ISO 9241 [25]). Ultimately, modelling event causality is complex, moreover, the effort may not be warranted for non-safety critical system, so three approaches are offered. First is to use the influencing factors as a paper-based 'tool for thought'. Second, the factors are implemented as a hypertext that can be traversed to explore contextual issues that may lead to errors and hence to generic requirements to deal with such problems. However, many of these variables interact, e.g. high stress increases fatigue. Finally as many combinations of influencing factors are possible and each domain requires a particular model, hence we provide a general modelling tool that can be instantiated with domain specific information. The tool allows influencing factors to be entered as a rating on a five point scale (e.g. high task low then calculates the event probability from the ratings. The combination of factors and ratings are user controlled. he factors described in figure 9 may be entered into the tool with simple weightings to perform sensitivity analyses. A set of default formulae for inter-factor weights are provided, but the choice depends on the user's knowledge of the domain. The tool can indicate that errors are more probable given a used defined subset of influencing factors, but the type of exception is difficult to predict, i.e. a mistake may be more likely but whether this is manifest as a event being omitted or in the wrong order is unpredictable. Where more reliable predictions can be made new 'alternative path' rules (see above) are added to the system. The tool is configurable so more validation rules can be added so the system can evolve with increasing knowledge of the domain. The current rules provide a baseline set that recommend generic requirements for certain types of agent e.g. untrained novices need context sensitive help and undo facilities, whereas experts require short cuts and ability to build macros. The influencing factors may be used as agent and use case properties and validated using the frames described in section 3.4. 3.7 Scenario Validation CREWS-SAVRE is loosely-coupled with Rational's RequisitePro requirements management tool to enable scenario-based validation of requirements stored in RequisitePro's data base. CREWS-SAVRE either presents each scenario to the user alongside the requirements document, to enable user-led walkthrough and validation of system requirements, or it enables semi-automatic validation of requirements through the application of pattern matching algorithms to each scenario. Each approach is examined in turn. Figure shows a user-led walkthrough of part of one scenario for the dealing system use case, and the RequisitePro requirements document being validated. The left-hand side of the screen shows a normal course scenario as a sequence of events. On the right-hand side are alternative courses and generic exceptions generated automatically from the requirements engineer's earlier selection of exception types. For each selected event the tool advises the requirements engineer to decide whether each alternative course is (a) relevant, and (b) handled in the requirements specification. If the user decides that the alternative course is relevant but not handled in the requirements specification, s/he can retrieve from CREWS-SAVRE one or more candidate generic requirements to instantiate and add to the requirements document. Each exception type in CREWS-SAVRE's data base is linked to one or more generic requirements which describe solutions to mitigate or avoid the exception. Thus, CREWS-SAVRE provides specific advice during user-led scenario walkthroughs. Figure 9 Influencing Factors for Exceptions and their interrelationships. Urfr.h#".r Gvtu#vt I'v+r X'.xfyhpr Q.rqvp#hivyv#' W'y"r Wh.vr#' H"y#v#h+xvt F'#yrqtr# @'fr.vrpr Brr.hy6f#v#"qr U.hvvt Q".B"vqhpr Ghpx's rp'".htrr# Q".qv+pvfyvr Q".xqr+vt CvqqrA"p#v'hyv#' Hv++vtsrrqihpx Hv+yrhqvtsrrqihpx Q".f.rqvp#hivyv#' Q'#r.+ffy'shvy".r 9v+pr. Ch.q#h.rshvy".r Ah#vt"r 8'tv#vo/oor shp#'.+ vp.rh+r 's Consider the example shown in Figure 10. The user is exploring normal course event 90, the start of the communication action 'the dealer enters deal information into the dealer- system' ,(shown in (A) in Figure 10s), and the alternative course GAc6, the information is suspect (B). The user, having browsed candidate generic requirements (C), copies and pastes the requirement 'the system shall cross-reference the information with other information sources to ensure its integrity' into RequisitePro's requirements document (D). This figure also shows the current hierarchical structure of the requirements held in RequisitePro's data base (E). Figure Validation Frames for Adding Generic Requirements The second approach automatically cross-checks a requirements document and a scenario using a collection of patterns which encapsulate 'good' socio-technical system design and requirements specification. To operationalise this, the CREWS-SAVRE tool applies one or more validation frames to each event or event pattern in a user-selected scenario to determine missing or incorrect system requirements. Each validation frames specifies a pattern of actions, events and system requirements that extend KAOS goal patterns [12] describing the impact of different goal types on the set of possible system behaviours. A validation frame has two parts. The first defines the situation, that is, the pattern of events and actions expressed in the form <identifier, action-type, agent-types> where agent-types are involved in the action. Each event is expressed as <identifier, event-type, action- identifier> where event type defines whether the event starts or ends the action. The second part of the frame defines generic requirements needed to handle the event/action pattern. The frames start from the PS055 standard [35] and type each requirement as a performance, usability, interface, operational, timing, resource, verification, acceptance testing, documentation, security, portability, quality, reliability, maintainability or safety requirement. Hence, automatic requirements-scenario cross-checking is possible using patterns of event, agent and action types in the scenario and requirement types in the requirements document. An example validation frame is: validation-frame {detect periods of system inactivity} situation: agent(agC,machine) AND not consecutive(agC,evA,evB); requirements: requirement(performance, optional, link); requirement(function <time-out/re-send>, optional, link); end-validation-frame This frame detects the absence of a reply event after a set time period from a human agent (implicitly) and signals the requirement to ask for resend or set a time out. An instantiation of this is the requirement to request a price to be entered by the dealer within seconds of accessing the 'prepare quote' option. This exception deals with inbound event delays, so if the time is longer than a preset limit (i.e. the system is not used for that period), the tool recommends reusable generic requirements, for example to warn the user and log-out the user after a certain period of time. Validation frames for alternative course events provide generic requirement to handle an event exception (ev) linked to an event 'ev' as follows: validation-frame situation: agent(agA,machine) AND requirements: requirement(functional, mandatory, link), system shall check for data entry mistakes), system shall restrict possible data entry); end-validation-frame Such attention failures are possible when entering dealer information. The tool proposes generic requirements, and the requirements engineer chooses requirement GR1 "the system shall check for data entry mistakes". In Figure 11, the requirements engineer is again examining event S90 [the dealer enters the deal information into the dealer-system] (A) and the causal influencing factor [agent pays poor attention to detail] (B). Again the above validation frame detects the need for functional requirements to handle the alternatives linked to event S90. As a result, the user is able to add new requirements to the requirements document to handle this alternative course. Figure Validation Frames for Exceptions Finally the tool uses validation frames for applying OSM-specific generic requirements in a similar manner. In section 3.4, high level generic requirements were recruited as design rationale trade offs, whereas when frames are used to detect OSM specific patterns in scenarios, more detailed requirements can be indicated. The computerised dealing system is an instantiation of the object supplying, object messaging and agent-object control system models. Consider a validation frame linked to the 'send' communication action in the object messaging system model. The situation (scenario) specifies an event which starts a communication action that involves a machine agent and which matches one of the actions in the object messaging OSM (i.e. sending or returning messages). For this situation, the validation frame identifies at least 5 generic system requirements: validation-frame comment: generic requirements for send action, object messaging OSM situation: [agent(agA,machine) requirements: requirement(functional, mandatory, link), generic-requirement(GOM1, system shall support identification/retrieval of receiver agent's address); system shall enable a user to enter the receiver agent's system shall enable a user to enter the content of a message); system shall enable a user to send a composed message to the receiver agent); system shall maintain a sender-agent log of all messages which are sent from and received by the sender-agent); end-validation-frame For example, the validation frame is applicable to event which starts the action 'buyer requests quote from dealer'. Each of the generic requirements is applicable to computerised support for this action, whether or not the action is undertaken by telephone or e-mail, for example retrieving the receiver agent's telephone number or e-mail address, and maintaining a sender log of all telephone calls or e-mail messages. Such generic requirements can be instantiated and added to the requirement specification. 4. Discussion The contributions to RE we have reported in this paper are threefold. First, extensive reuse of requirements knowledge is empowered via decision models and generic requirements. Secondly a means of semi-automatic requirements validation is provided via frames. Frames extend type checking by recognising patterns of agents' behaviour to which appropriate validation questions, and possible design solutions, may be applied. Third, we have described the use of scenarios as test pathway simulations, with novel tool support for semi-automatic scenario generations. The current status of development is that the scenario generator-validator tool has been implemented and industrial trials are commencing. Clearly coverage in terms of the number of validation frames and generic requirements contained in the tool database is a key issue effecting the utility of our approach. Our approach is eclectic and depends on knowledge in literature such as from ergonomics, human resource management and user interface design. The contribution we have made to implement and integrate this knowledge in an extensible architecture. The advice currently contained retrieved by the validation frames provides requirements knowledge at a summary level. Although this may be criticised as lacking sufficient detail initial industrial reaction to the tool has been encouraging, in particular the value of raising design issues which are not addressed by current methods, e.g. DSDM [15]. So far we have demonstrated proof of concept in terms of operation. This will be followed by further testing of utility within a small scale, but industrially realistic application. In many senses the strength of the method we have proposed lies in integration of previous ideas. We have brought concepts from safety critical system assessment [24], [42] to bear on requirements analysis, and integrated these with scenario based approaches to RE. We acknowledge the heritage of the Inquiry Cycle [39]; however, our research has contributed an advanced method and support tool that give more comprehensive guidance for solving problems. Specification of requirements to deal explicitly with the implications of human error is a novel contribution where we have broken ground beyond the previous approaches [30], [16]. Furthermore, the influencing factors that bear on causes for potential error are useful heuristics to stimulate debate about many higher level requirements issues, such as task and workplace design. However, we acknowledge it is difficult to provide prescriptive guidance from such heuristics. While some may contend that formalising analytic heuristics can not capture the wealth of possible causes of error in different domains, we answer that some heuristics are better than none and point out that the method is incremental and grows by experience. Failure to formalise knowledge can only hinder RE. Parts of the scenario based method reported in this paper are related to the enterprise modelling approach of Yu and Mylopoulos [54] and Chung [7]. They create models of the system and its immediate environment using similar semantics for tracing of dependencies between agents, the goals and tasks with limited reasoning support for trade-offs between functional requirements and non-functional requirements (referred to as soft goals). However the i* method does not contain detailed event dependency analysis such as we have reported. Scenarios have been used for assessing the impact of technical systems by several authors [6], [30], [16]. However, these reports give little prescriptive guidance for analysis, so the practitioner is left with examples and case studies from which general lessons have to be extracted. For instance, the ORDIT method [16] gives limited heuristics that advise checking agent role allocations, but these fall far short of the comprehensive guidance we have proposed. Dependencies between systems and their environment have been analysed in detail by Jackson and Zave [28] who point out that input events impose obligations on a required system. They propose a formalism for modelling such dependencies. Formal modelling is applicable to the class of systems they implicitly analyse, e.g. real time and safety critical applications, but it is less clear how such models can deal with the uncertainties of human behaviour. To deal with uncertainty in human computer interaction, we believe our scenario based approach is more appropriate as it focuses on eliciting requirements to repair problems caused by unreliable human behaviour. Another approach is the KAOS specification language and its associated GRAIL tool [32], [33] a formal modelling that refines goal-oriented requirements into constraint based specifications. Van Lamsweerde et al [33] have also adopted problems and obstacles from the Inquiry cycle [39]; furthermore, they have also employed failure concepts from the safety critical literature in a similar manner to CREWS-SAVRE. Their approach is anchored in goal-led requirements refinement and does not use scenarios explicitly. In contrast, CREWS-SAVRE covers a wider range of issues in RE than KAOS but with less formal rigour, representing a trade-off in RE between modelling effort, coverage and formal reasoning. So far the method has only partially dealt with non functional requirements. Scenarios could be expressed in more quantifiable terms, for instance by the Goal- Question-Metric approach of Basili et al. [4], or by Boehm's [5] quality-property model. Scenarios in this sense will contain more contextual information to represent rich pictures of the system and its environment [29]. The description could be structured to include information related to the NF goal being investigated and metrics for benchmark testing achievement of the goal. Validation frames may be extended for assessing such rich picture scenarios for non functional and functional requirements. For instance, each inbound/ outbound event that involves a human agent will be mediated by a user interface. Usability criteria could be attached to this event pattern with design guidelines e.g. ISO 9241 [25]. Performance requirements could be assessed by checking the volume and temporal distribution of events against system requirements. Elaborating the scenario based approach to cover non functional requirements is part of our ongoing research [52]. In spite of the advances that scenario based RE may offer, we have still to demonstrate its effectiveness in practice. There is evidence that the approach is effective in empirical studies of earlier versions of the method which did use scenarios but without the support tool [50]. Further validation with industrial case studies is in progress. Acknowledgements This research has been funded by the European Commission ESPRIT 21903 long term research project 'CREWS' - Co-operative Requirements Engineering With Scenarios. The project partners include RWTH-Aachen (project co-ordinator), City University, London, University of Paris I, France, FUNDP, University of Namur, Belgium. --R Object Oriented Analysis Dynamic Systems Development Method (DSDM) Version 3.0 Analysis Patterns: Reusable Object Models Design Patterns: Elements of reusable object-oriented software ISO 9241 Human Factors Engineering and Design Human Computer Interface Design "Scenario-based Analysis of Non-Functional Requirements" --TR --CTR Lin Liu , Eric Yu, Designing information systems in social context: a goal and scenario modelling approach, Information Systems, v.29 n.2, p.187-203, April 2004 Norbert Seyff , Paul Grunbacher , Neil Maiden , Amit Tosar, Requirements Engineering Tools Go Mobile, Proceedings of the 26th International Conference on Software Engineering, p.713-714, May 23-28, 2004 Stan Jarzabek , Wai Chun Ong , Hongyu Zhang, Handling variant requirements in domain modeling, Journal of Systems and Software, v.68 n.3, p.171-182, 15 December John A. van der Poll , Paula Kotz , Ahmed Seffah , Thiruvengadam Radhakrishnan , Asmaa Alsumait, Combining UCMs and formal methods for representing and checking the validity of scenarios as user requirements, Proceedings of the annual research conference of the South African institute of computer scientists and information technologists on Enablement through technology, p.59-68, September 17-19, Zhang , Dan Xie , Wei Zou, Viewing use cases as active objects, ACM SIGSOFT Software Engineering Notes, v.26 n.2, March 2001 Sebastian Uchitel , Greg Brunet , Marsha Chechik, Behaviour Model Synthesis from Properties and Scenarios, Proceedings of the 29th International Conference on Software Engineering, p.34-43, May 20-26, 2007 Emmanuel Letier , Jeff Kramer , Jeff Magee , Sebastian Uchitel, Monitoring and control in scenario-based requirements analysis, Proceedings of the 27th international conference on Software engineering, May 15-21, 2005, St. Louis, MO, USA Dalal Alrajeh , Alessandra Russo , Sebastian Uchitel, Inferring operational requirements from scenarios and goal models using inductive learning, Proceedings of the 2006 international workshop on Scenarios and state machines: models, algorithms, and tools, May 27-27, 2006, Shanghai, China Nico Lassing , Daan Rijsenbrij , Hans van Vliet, How well can we predict changes at architecture design time?, Journal of Systems and Software, v.65 n.2, p.141-153, 15 February Giuseppe Della Penna , Benedetto Intrigila , Anna Rita Laurenzi , Sergio Orefice, An XML environment for scenario based requirements engineering, Journal of Systems and Software, v.79 n.3, p.379-403, March 2006 Julia Galliers , Alistair Sutcliffe , Shailey Minocha, An impact analysis method for safety-critical user interface design, ACM Transactions on Computer-Human Interaction (TOCHI), v.6 n.4, p.341-369, Dec. 1999 PerOlof Bengtsson , Nico Lassing , Jan Bosch , Hans van Vliet, Architecture-level modifiability analysis (ALMA), Journal of Systems and Software, v.69 n.1-2, p.129-147, 01 January 2004 Sascha Konrad , Betty H. C. Cheng , Laura A. Campbell, Object Analysis Patterns for Embedded Systems, IEEE Transactions on Software Engineering, v.30 n.12, p.970-992, December 2004 Idris Hsi, Measuring the conceptual fitness of an application in a computing ecosystem, Proceedings of the 2004 ACM workshop on Interdisciplinary software engineering research, November 05-05, 2004, Newport Beach, CA, USA Andreas Gregoriades , Alistair Sutcliffe, Scenario-Based Assessment of Nonfunctional Requirements, IEEE Transactions on Software Engineering, v.31 n.5, p.392-409, May 2005 Colette Rolland , Naveen Prakash, From conceptual modelling to requirements engineering, Annals of Software Engineering, v.10 n.1-4, p.151-176, 2000 Alistair Sutcliffe, On the effective use and reuse of HCI knowledge, ACM Transactions on Computer-Human Interaction (TOCHI), v.7 n.2, p.197-221, June 2000 Axel van Lamsweerde , Emmanuel Letier, Handling Obstacles in Goal-Oriented Requirements Engineering, IEEE Transactions on Software Engineering, v.26 n.10, p.978-1005, October 2000 Robb Klashner , Sameh Sabet, A DSS Design Model for complex problems: Lessons from mission critical infrastructure, Decision Support Systems, v.43 n.3, p.990-1013, April, 2007 Marcos Andr Gonalves , Edward A. Fox , Layne T. Watson , Neill A. Kipp, Streams, structures, spaces, scenarios, societies (5s): A formal model for digital libraries, ACM Transactions on Information Systems (TOIS), v.22 n.2, p.270-312, April 2004 N. Robinson , Suzanne D. Pawlowski , Vecheslav Volkov, Requirements interaction management, ACM Computing Surveys (CSUR), v.35 n.2, p.132-190, June

scenarios;use cases;reuse;requirements engineering;scenario generation;exception types;patterns

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The Use of Cooperation Scenarios in the Design and Evaluation of a CSCW System.

AbstractDesign and evaluation of groupware systems raise questions which do not have to be addressed in the context of single user systems. The designer has to take into account not only the interaction of a single user with the computer, but also the computer-supported interaction of several users with each other. In this article we describe the use of cooperation scenarios in the design and evaluation of an innovative access control system for a concrete groupware application developed in the POLITeam project. We have used informal textual scenarios to capture a rich description of the particularities of access to cooperatively used documents in three different organizations. Based on these scenarios, we have developed an access control system, which not only allows specification of access rights in advance but also allows involvement of third persons at the actual time of access, using negotiation and notification mechanisms. We describe our evaluation strategy which again employs the cooperation scenarios developed in the empirical phase. After relating our approach to other work, we summarize and discuss our experiences and the advantages (and disadvantages) of using scenarios for the design and evaluation of Computer Systems Cooperative Work (CSCW) systems. Finally, we give a brief outlook on future work.

Introduction COMPUTER SUPPORTED COOPERATIVE WORK (CSCW) is an interdisciplinary field of research, dealing with cooperative work supported by computer systems (for an overview see e.g. [6]). Computer science naturally is a major contributor. However, psychology, sociology, and other disciplines are involved, as well. The developers of CSCW systems (groupware) not only have to design for single users interacting with the system but also for groups interacting via the system. This collaborative perspective raises or amplifies, for instance, issues like privacy and access control ([9, 23]), conflicts ([27]), awareness ([4, 7]), and tailorability ([21, 22]), which we discuss briefly in the following in order to give the reader an overview of some of the challenges in groupware design. As early as 1986, in the context of the first CSCW conference, Greif and Sarin ([9]) stated that the access control mechanisms and concepts used in operating systems of that time were not flexible enough to express access policies for group interaction. They suggested the development of more sophisticated controls taking into account additional factors like user roles, access rights concerning abstract operations other than read and write (e.g sharing operations), and specific object-user-relationships (e.g. the current or past users of an object). Shen and Dewan ([23]) discuss the problem of access rights for collaborative work in the context of the multi-user editor SUITE. They stress the dynamic nature of collaborative work and thus state the support of multiple, dynamic user roles and the need for easy specification of access rights as important requirements for access control in groupware. Furthermore, they introduce specific collaboration rights for operations whose results can affect other users. The fact that groupware functionality can affect multiple users gives rise to a strong potential for conflict concerning the configuration and use of such systems ([27]). Groupware systems can distinctively change the division of labor in an organization, obliterate jobs, and open new opportunities for communication between employees (a development which might not be in the interest of management). Furthermore, supporting communication and coordination with computers opens new ways of controlling and monitoring work. Often small things like the publicly visible creation date of a document are cause for vehement negative reactions of groupware users. While groupware systems allow for more control over work processes on one hand, on the other hand important context information can be obliterated by cooperating e.g. via a shared document work space. Even small clues like seeing the overflowing (physical) inbox of a colleague or hearing the voice of a somebody in the hallway might have formerly been helpful in deciding questions like whom to send which document or when to ask for the completion of some piece of work. When supporting cooperative work with computers, this raises the issue of awareness (see e.g. [4]). Groupware users sometimes need to be aware of what is going on in the systems and what other users are currently doing or have done in the past. Designing mechanisms to provide awareness in groupware is nontrivial because one has to walk a tightrope concerning the conflicting goals of privacy versus the need for awareness and information overflow versus lack of awareness. Fuchs et al. ([7]) present as one solution to this problem a model for an awareness mechanism which is characterized by a high degree of tailorability, i.e. it can be adapted to different and dynamically changing needs and preferences of individuals, groups, and organizations. While tailorability (or adaptability) is also an issue outside CSCW, the complexity, dynamics and diversity of cooperative work increase its importance in this field beyond the configuration of awareness mechanisms (see e.g. [21, 22]). Furthermore, the fact that groupware functionality has the potential to affect multiple users raises the question of who is allowed to tailor this functionality and how one can explore and try out tailored functionality without disturbing other users. Access control, potential for conflicts, awareness, and tailorability are some examples (and by no means all) of the issues which arise in the design of CSCW functionality. However, when extending the use of computers from the support of single users to the support of groups, not only the functionality but also the development processes have to be rethought. [10] gives an extensive overview of the novel challenges groupware designers have to face. Among others, he identifies the need for almost unanimous acceptance of groupware in order to achieve a critical mass of users and a well balanced benefit profile (e.g. subordinates gain as much from the system as managers) as key factors to the success of groupware. Thus, groupware developers have to know much more about the context of use of the prospective application than developers of single user applications, who do not have to pay as much attention to group-related aspects like, for instance, trust, awareness of each other, negotiations, social dynamics, power structures, and work processes. Trust, awareness, and negotiation are aspects which proved highly relevant in the design case described later on. [10] also points out the difficulties in evaluating groupware. The interaction of a single user with a system is much easier to evaluate in a laboratory setting than group processes. Apart from the sheer logistical problems of getting even a small-sized group in the same lab at the same time or the nightmare of installing a prototypical application in an organization, the necessary duration of the evaluation is a major problem, because "group interactions unfold over days or weeks" ([10], p. 101). In the following we want to concentrate on the two design process related problems identified above: capturing of rich contextual requirements and evaluation support. In the POLITeam project (see [17]), we employ textual scenarios drawn from field studies, interviews, and workshops to inform the design processes and support the evaluation prior to field tests and large scale workshops. We call our scenarios cooperation scenarios, as they not only capture the work and interaction of single users with the system but also the group and organizational context and its work practices. Within the POLITeam project a groupware application for a German federal ministry and selected ministries of a state government and the concurrent engineering division of a car producer is developed in an evolutionary and participative way. The first system version was generated by configuring the commercial product LINKWORKS by Digital. Based on the experiences gained by introducing the first system version in three different fields of application, we develop advanced versions of the system. The functionality mainly consists of an electronic circulation folder, shared workspaces, and an event notification service. In this article we describe the use of cooperation scenarios in the design of the access control system in POLITeam. Specifically, we want to point out the value of cooperation scenarios for the creation of novel CSCW-functionality, in this case the integration of a traditional anticipative access control system with negotiation (computer mediated decision making, see also [27]) and notification services. In section two we describe the concept of cooperation scenarios in more detail, comparing it with other types of scenarios in the related literature. Section three introduces the access control design problem and presents three cooperation scenarios which have notably guided and motivated the new design. Section four contains the analysis of the scenarios and briefly outlines the resulting implementation. Section five describes the use of scenarios in three consecutive evaluation steps. Section six relates our approach to other work. Section seven summarizes and discusses the value of cooperation scenarios. Finally, section eight suggests future research efforts. Cooperation Scenarios for the design and evaluation of CSCW systems [3] identifies several different roles which scenarios can play in the development process of software. We employ scenarios in the CSCW design process in three roles: as a tool for the first (informal) requirements analysis, for communication support (user-designer and designer-designer) during validation, and for evaluation support. Figure 1 shows how the scenario supported steps are positioned in the overall design process. Note that the process is cyclical and that we believe the value of cooperation scenarios lies foremost in the first cycles when attempting to identify or create innovative functionality which supports cooperative work. Later cycles rely on more formal methods and models. The form of cooperation scenarios Scenarios can take many different forms. [18] describes two extreme positions concerning the scope of scenarios. The first one "sees a scenario as an external description of what a system does" (p. 21), while the second one looks "at the use process as situated in a larger context" (same page). According to these different positions, scenarios can take the form of exact protocols or formal sequence descriptions on one hand, and rather broad, mostly textual descriptions also covering contextual aspects which are only loosely related (however, relevant) to system design on the other hand (also see [3]). For our purposes we need a form of scenario, which allows designers to capture a broad range of (eventually) unanticipated, contextual information. Thus, our cooperation scenarios are based on informal textual descriptions of work practices, including the motivation and goals behind cooperation. The textual form also facilitates the discussion between users and designers (role of communication support) and among designers. This first representation can later be augmented with alternatives, especially visual techniques. Workshops Work place visits Requirements (informal textual scenarios) Design functionality Evaluation without users e.g. role-based analysis of scenarios yielding a benefit profile Evaluation with users Presentation and discussion of new functionality based on scenarios Re-design Re-design Validation of scenarios Use Cyclical design in the POLITeam project Scenario based process steps Figure 1: Scenario support in the POLITeam design process Building cooperation scenarios The information captured in cooperation scenarios is gathered in the POLITeam project through extensive field studies, involving semi-structured interviews and work place visits. While these techniques are usually single user oriented, we attempt to identify the role a person plays in the organizations and the different cooperative tasks. The question of individual motivations and goals is especially important in the context of groupware development, as the product has to be acceptable to users in radically different roles with respect to the cooperative activity which is being supported by the system (e.g. managers and subordinates, see [10]). For this purpose, we use what we call continuously refined heuristic user selection schemes (see [25]). Having identified a cooperative activity, we attempt to identify the users involved in this activity. For the interviews and work place visits, we select users who play different roles in the cooperation. However, our initial selection of users is often not exhaustive and as we learn more about the processes in our target organizations, we identify additional users, who might play a different role in a more subtle, however relevant, way. Thus, we continuously refine the user selection. A methodologically important aspect of cooperation scenarios is the degree to which the cooperation in the respective field of application is already supported by computer systems. The three scenarios described in the next section range from full computer support to completely non-IT based activities which might be subject to future groupware support. The existing work practices are usually heavily influenced by the features (and problems, see scenario one) of the computer system already in place. It is essential to abstract from existing technology in order to understand the nature and especially the motivation and goals behind the cooperative activity. The main point of this article is to make a case for the value of scenarios as support for developing this understanding. Additionally, we will discuss the value of scenarios in the evaluation phase. Three example scenarios concerning access control in the context of collaborative activities When the first POLITeam version was introduced, it quickly became obvious that the users had severe problems with the traditional (matrix based) access control system (see e.g. [20]). They just did not use it. In order to understand the nature of these problems, we began to investigate the work practices revolving around access to collaboratively used documents. The three scenarios presented in this section were distilled from a series of interviews in two fields of application of the POLITeam project, an additional interview session in a newsletter's editorial office, and a workshop with users from the federal ministry involved in the POLITeam project. The workshop was also concerned with other questions (which are not relevant here) and is not to be confused with the feedback workshop described later on in this paper. The scenarios and their analysis are taken from [26], where they are used to motivate the novel aspects of the access control system. (The main points of [26] are technical aspects of the implementation of this system and thus it employs a more formal method than scenarios (petri nets) in order to describe and discuss implementation details and alternatives at the heart of the novel functionality. [26] does not discuss the design process as such). First scenario: a state representative body - keeping user passwords in sealed envelopes in a strongbox This scenario is based in the representative body of a North German state at the federal capital in Bonn. As participant in the POLITeam project, the body has been equipped with a groupware system, supporting internal cooperation as well as the cooperation with the state government in the state capital. The installation consists of 28 workstations. An important cooperative task being supported is the preparation of documents for the state's vote in the Bundesrat (the German assembly of state's representatives). The cooperative activity is rather time constrained, as decisions concerning the state's vote have to be coordinated, distributed, and validated within a very short time frame (about one week). In order to deal with unexpected absences due to illness or travel, the office has implemented a seemingly complex, but rather effective non-IT-based work practice around the POLITeam groupware system. Every user has to write his password on a piece of paper, which is placed in a sealed envelope. This envelope is locked away in the office strongbox. The keys to the strongbox are kept by two trusted persons: the system administrator and the department head) (in the following referred to as key-holders). If a person urgently needs to access a document on the "virtual desk" of an absentee, the person has to ask one of the key-holders to open the strongbox and release the envelope with the password. The envelope is opened and the "virtual desk" can then be accessed with the appropriate password. On one hand, this work practice is effective in the sense, that misuse or illegitimate access is very difficult because of the necessity to negotiate with the key-holder to release the envelope and because the broken seal of the envelope is an indicator that the virtual desktop was accessed. On the other hand, the granularity of the access granted by the system is rather coarse, since once the password is released, every possible action can be taken in the name of the absent person. Other documents than the one needed can be accessed or deleted; mails can be sent and received. Additionally, the eventual change of a user password is causing some organizational overhead, because the password not only has to be changed in the system, but also a new envelope has to be placed in the strongbox. Theoretically, a person could use several virtual desktops with different password to achieve a finer granularity. This solution, however, would severely decrease the effectiveness of the regular use of the system, since the users would have to log in and out of the system if they change from one task to another. Second scenario: a federal ministry - searching your colleagues' desktops under the watchful eyes of a trustworthy third person A different scenario was discovered in a department of a German federal ministry. The department also participates in POLITeam. The groupware application is used by 12 employees mainly in an operative section and the central typing office. The application supports the cooperative generation of documents, typically involving the head of the section, a member of the section and several typists. In the course of their work, the members of the section occasionally have to search for documents on other colleagues' virtual desktops, for instance, if they need a specific document, but do not know who is working on it right now. For this task the POLITeam base system provides a search tool, which allows users to search for and access all documents, which are not explicitly declared private. This tool, however, was severely constrained in the installed version of the groupware system, because in the initial requirement analysis, the users objected to the tool due to privacy considerations. The current version of the tool can only search the folder-hierarchy on one's own virtual desktop. Conventions have been established, which deal with the problem of searching for work outside one's own virtual desktop. Users are held to keep documents which are jointly created in shared workspaces. Each member of the section has a workspace whose access he shares with the head of the section and the typists. Thus, the two typists working for the section have a link to all shared workspaces on their desktop. This means that they can conduct a department-wide search for documents. If one of the workers in the section looks for a specific document outside his own desktop, he asks them to search for him and provide him with a link (a POLITeam concept for shared access to documents) to the found document. The advantage of this work practice is that the two typists are aware of every department-wide search and - as they are considered extremely trustworthy - nobody feels that his privacy is being invaded. The disadvantage is that access to documents still has to be anticipated, in order for these documents to be kept in shared workspaces. Furthermore, because the typists actually have to search for and copy or link the respective documents, this practice results in additional work for them, even though its sole point is that they are aware of the search and can intervene, if they doubt its appropriateness. Third scenario: newsletter editorial offices - smoothing cooperation with limited trespassing in private domains Another example for a successful work-practice from the physical world, which is very difficult to support in classical access control systems, was encountered in the editorial offices of a small newsletter, specialized on providing in-depth information about EU agricultural matters to interested parties. The seven editors are responsible for up to three EU countries each. Thus, in theory the work of the editors is non-collaborative. However, in practice there is lot of information which concerns transnational aspects, e.g. studies comparing agricultural performance of several EU countries or simply a newspaper article from one country which concerns another country. As a consequence, the editors heavily rely on being kept up to date by their colleagues. The supporting mechanism for this collaboration is a (physical) circulation system, which is based on a number of open post-boxes in the main hall of the suite of offices. The post- 6boxes are also used for distributing the normal mail to the editors. If an editor wants to share a document with his colleagues, he simply writes the initials of the respective persons on the document on drops it in one of the post-boxes. Now, sometimes the editor sharing the information decides afterwards that he needs the document himself again. Or he might tell one of his colleagues about the document over lunch, who might discover that this is exactly the information he has been looking for for weeks and immediately needs to see the document. The document in question, however, still resides in one of the post-boxes in the main hall. Thus, in order to speed up the process, the editor searches the post-boxes of his colleagues. However, he would not dare to remove any other documents or look too closely at private mail or faxes, because the post-boxes are located in the main hall and it is very likely that a colleague might pass by. Again, similar to the last scenario, the editor is respecting a social protocol. According to the interviews this custom is honored even by the last person leaving the office at night.] When trying to support this kind of work-practice in a computer system, one quickly reaches the limit of classical access control system. Analyzing and Using the Scenarios to Guide the Design Process We have used the three scenarios described above as the basis for the development of a new access control model. We found cooperation scenarios very useful at this point, because they revealed to us some problems with traditional access control systems in collaborative work practice which we had not seen before. In the following, we briefly describe three observations (taken from [26]) which indicate why traditional access control mechanisms do not support current collaborative work practice very well. These observations are based on the information captured in the three scenarios and guided the subsequent design process, in particular the integration of negotiation and notification mechanisms into the access control system. Observation 1: Trusted third persons play an important role in the three scenarios The role of a trusted third person is a central feature in the first two scenarios. In state representative office, the employees put the envelopes with their passwords in the strongbox, well knowing that there are others who have the keys. The key-holders are trusted not to abuse their powers. Similar in the federal ministry the typists are considered trustworthy by their colleagues. Thus, they are allowed to search for documents in restricted spaces. Trusted third persons obviously play an important role when dealing with unexpected absences. This is exactly the point where classical access control system fail. These systems only deal with two roles, the one who specifies the access rules and the one who tries to access a document. Access policies, e.g. in form of access lists or matrices, have to be specified in advance, precisely denying or permitting access to an object for certain persons or groups. The problem in the two scenarios is, that the users would like unexpected accesses not to be decided upon in advance, but by a trusted third person in the context of the actual access situation, e.g. by the typists or the In theory these scenarios could be supported by a traditional access control system by just granting full access to the trusted person in advance. This solution, however, has the consequence that the trusted person would not only be burdened with the decision concerning access but also has to do the actual work, i.e. accessing and copying the document for the person requesting the access as in the scenario from the federal ministry. If unanticipated accesses are not exceptional, this solution is not acceptable for the trusted persons. Thus, it makes sense to look for a more efficient implementation of these requirements within the access control system. Observation 2: Awareness can be used to control access Awareness plays an important role in all three scenarios. In the state representative office, the passwords are sealed in envelopes. A broken seal indicates that the password was used. Thus, there is a degree of awareness about accesses to one's desktop. In the federal ministry, access to restricted spaces has to be made aware to the typists. This awareness ensures that every access is carefully considered because one can be held accountable for every action. The situation in the editorial offices is rather similar. The position of the post-boxes in the main hall provides a degree of awareness, which ensures that other editors' boxes are only accessed, if there is a justifiable reason. Awareness plays an important role by supporting and enforcing social protocols or conventions. Again this is the point where classical access control systems fail. They only allow for the options yes and no, not yes, but I want to know about it or yes, but I want the typists to know about it. The POLITeam base system includes functionality, which allows users to register interest in changes to specific documents. This functionality, however, is much too coarse in that it does not differentiate between different kinds of accesses. Additionally, it can only be specified for one document at a time, which induces a lot of overhead when trying to use the functionality in real work practice. Observation 3: Access can be subject to negotiation at the time of access Looking at the work practices described in the scenarios, it is obvious that there is a great need for negotiation. If it is, for any reason, not possible to precisely specify access rights in advance, the anticipated specification is replaced by negotiate [sic] during the actual situation of use, either with a trusted third person (see first subsection) or the owner of the document himself. In our scenarios, the necessity for negotiation is a direct consequence of the application of trust in the respective organization. In the scenario taken from the state representative body, one has to negotiate with one of the key-holders about the release of the password. The key-holder has the opportunity to grant or deny access according to his understanding and motives in the current context. In the federal ministry one has to negotiate with one of the typists in order to search for certain documents. Implementation Based on this analysis, we have developed a new access control model, which integrates awareness and negotiation services with a traditional, anticipative access control system. When specifying access rights for a document or a shared folder, the users not only have the opportunity to allow or deny access in advance. They have two additional options: notification and negotiation. The former allows access, but notifies a single user or a group about the access. The latter only allows access, if a specified single user or group agrees. For a detailed discussion of the implementation see [26]. Scenario supported evaluation So far, we have described the use of cooperation scenarios for the purpose of capturing requirements which stem from cooperative work practice and guide the design of new CSCW functionality. However, we also found scenarios very useful during the evaluation phase. As mentioned in the introduction, evaluating groupware is rather difficult, due to the many factors involved in success or failure of a product. A full-fledged test with groups of users - preferably at their work place - is impractical before the product has reached a certain maturity (especially stability). If groups of users are invited to the lab, they leave a lot of important context factors behind and are subjected to several new ones. Our strategy is to capture as much of the relevant context as possible in cooperation scenarios, especially the motivation, the goals, and also the workload of the different participating roles. We try to evaluate our designs against this information as early as possible in the process. Specifically, we use scenarios for evaluation in three stages of the design process (see figure 1): evaluation of scenario validity, theoretical evaluation of system design, and the practical evaluation in user workshops. Evaluation of scenario validity When we transcribed the scenarios presented in [26], we found that there were some differences in the interpretation of relevant details. This was despite the fact that both authors had attended the (first) workshop (which had been protocolled) and both authors conducted the relevant interviews together (which had been audio-taped). A good example for a significant misinterpretation was discovered when writing up and discussing scenario two. During the workshop, one of the attendees from the federal ministry was talking about searching other peoples desktops and the importance of other, trusted persons knowing about this search. She was referring to the "virtual" desktops of her colleagues, which was clear to everybody knowing that she had her office three stories above the colleagues she was talking about. However, the author writing up the workshop protocol did not know that and assumed she was talking about "physical" desktops. The mistake was finally discovered, when the other author reviewed the scenario. We believe that misinterpretations like this are a major problem in design processes with user participation. Common misconceptions can be cemented into what we call "project folklore", i.e. things that all designers believe in and which even are passed on to new project members as "facts". Our cooperation scenarios rely on anecdotal evidence which can be severely distorted depending on who writes it up and what prior knowledge this person has about the field of application in question. Thus, we found it very helpful, to exchange cooperation scenarios among designers after write-up and critically compare and discuss them, specifically asking questions like "I did not get this detail. Who said that and are you sure, you've interpreted it correctly?'' As a last measure, we sometimes even asked the interviewees later on, whether this or that detail matched reality. For this purpose a feedback workshop with all end users involved is quite useful. Theoretical evaluation of system design As mentioned before, we tried to capture not only the cooperative activity but also the motivation and goals behind individual contributions to the activity. Another important aspect is the workload of a specific role in the scenario. Having produced an early stage of the new design, we can already employ scenarios for a role-oriented benefit evaluation. We "insert" the new design, i.e. the new CSCW functionality, in the scenario and analyze for each role in the cooperative activity how individual parts of the task change. Take, for instance, the first scenario. In the old scenario, the system administrator or department head had to be asked to open the strongbox and extract the password from the envelope. In the new scenario, the documents are protected employing the negotiation service. How does the new scenario impact upon the individual parts of the tasks? Table 1 shows a role-oriented analysis of the old and new scenario one: Role Old scenario New scenario with negotiation service support User requesting access . Has to contact system administrator. . Has to wait for system administrator's decision . Has to log in under data owner's account . Has to copy the relevant documents to his own account . Can be held responsible for any additional changes to data owner's account . Has to wait for system administrators decision (System administrator or any other trusted person is contacted automatically) System administrator (department head is omitted here for simplicity) . Has to make decision . Has to open safe . Has to extract password from envelope . Has to communicate password to user . Has to change password later on . Has to create new envelope with new password . Has to answer the automatically generated request. Data owner . Does not know, what other data has been compromised . Has to submit his password to the system administrator in advance . Has to specify access rights and negotiation service in advance (can be a lot of work in dynamic work environments) Table 1: Role-oriented analysis of the new, technology-enhanced scenario one Role-oriented analysis of technology-enhanced scenarios can show, who benefits from the introduction of new technology and who has to carry additional burden (compare [10]). The analysis shown in table 1 indicates that the user requesting access and the system administrator benefit from the new functionality, while the data owner has additional work in specifying the configuration of access rights and negotiation services. While role-oriented analysis gives some hints concerning the possible acceptance of a system, there are still a lot of factors which might be relevant, but do not surface in the analysis. For example, it might be a huge benefit for the data owner for personal reasons, if he can specify another person than the system administrator as decision maker. Or there might be certain documents on a persons desktop which have to be excluded by law from the scheme (e.g. medical documents), which is possible with the fine-grained configuration of the negotiation service. These examples show the importance of capturing context in cooperation scenarios for theoretical evaluation. Practical evaluation of system design in workshops While theoretical analysis might yield some useful hints concerning possible problems with the groupware system, the almost canonical unreliability of any requirements analysis process indicates the need for user feedback as early as possible in the design process. For this purpose we conducted a feedback workshop, during which we presented an early (rather unstable) version of our prototype to end users (the feedback workshop is not to be confused with the earlier one). 11 end users from one of our fields of application (the federal ministry) participated in the workshop, among them the head of department. Seven designers were present; 3 from the University of Bonn and 4 from GMD (St. Augustin, Germany) where the workshop took place. The project members from GMD were not directly involved in the design of the access control system presented here, but contributed to the discussion during the workshop. The main goal of the presentation of the prototype was to help the users in understanding the new functionality achieved by our integration of traditional "yes or no" access control with negotiation and notification services. Based on this understanding we wanted to know from the users how they believed the system could be used to support (or change) their work practices. Implementation details (e.g. the user interface) of the prototype were also discussed, but are not in the scope of this paper, as are the usability tests we ran (not during the workshop) to evaluate the interfaces for configuring the negotiation and notification services. The presentation during the workshop was based on a scenario drawn from the federal ministry concerning the deletion of an address list in a common workspace. (not one of the scenarios described before). Together with student volunteers we "played" the technology enhanced scenario from a script, which we had ensured did not cause the prototype any trouble. The screens of two workstations were fed into beamers, so all participants could see the details. The workshop, especially the discussion and contribution of the end users, was protocolled. The real world cooperation scenario captured the imagination of the workshop attendees and helped us, first to discover some flaws in the base scenario and secondly to discuss the design of the new functionality (the users wanted more high-level and powerful configuration mechanisms for the negotiation service and the basic access rights, which supports the result from the theoretical analysis that the additional workload for the data owner might be a problem). Concerning the central question of how the new functionality might support their work practices, the users voiced rather diverse opinions. Interestingly, the head of the department said that if all responsibilities (e.g. for deleting address lists) were assigned "correctly", negotiation during the time of access would not be necessary. Furthermore, a high level of individualized (as opposed to standardized) procedures would actually hinder cooperation, because single contributions would not fit together. Both points were rather vehemently opposed by his subordinates, who stated that there were not one "correct" way to assign the responsibilities, but that in their everyday work practice many things were in flux and could not be cast into fixed and standardized structures. From each point of view, the statements are justified. The head of department wants to have clear (at best static) picture of what is going on in his department, while the subordinates have to actually deal with the surprises and non-standard situations in everyday work. This exchange of views exemplifies the group-related aspects which were mentioned in the introduction and which CSCW design has to take into account. The cooperation scenario we employed as basis for the presentation, helped in supporting the understanding of the end users and the subsequent discussion. Related work The approach presented here mainly builds on work which is concerned with putting the end user and his or her work in the focus of attention during software development. This perspective is due to the nature and the current state of CSCW. The discussion in this field is still in an exploratory phase. It is concerned with the question of what functionality groupware systems should provide (e.g. notification mechanisms) and what the benefits of these systems for individual users, groups of users and organizations really are. This reflects on the design methodologies which are used to support groupware design processes. While more traditional software design (see e.g. [24]) proceeds rather quickly from the initial requirements analysis and specification phase to formally verifiable application models and perhaps executable specifications, the designer of a groupware system has to spend much more time in the early phases of development, envisioning how the new functionality might change current work practice in an organization. Similarly, Hammer and Champy ([11]) postulate the "enabling role of information technology" (p. 83). Organizational and technical changes are intertwined and cannot be viewed or implemented separately. The introduction of a groupware system should not cement existing work practice in program code, but allow for new forms of collaboration which perhaps better support the organizational goals. Putting the end user and his or her work in the focus of attention is not a new concern in system development. As computers enter more domains of private and working life, novel uses of information technology are pioneered. This development necessitates design methodologies which put the use and the user of the system in the center of the development efforts. This does not mean that traditional, more formal techniques are obsolete. To the contrary, as system become more complex, formal methods are indispensable. However, they have to be complemented with a sound understanding of the current work practice and the possible future use of the system from the end users' perspective. The use-case methodology by Jacobsen et al. ([15]) is an example for the incorporation of the end user's perspective into system design. Use cases are descriptions of the interaction of an actor external to the system (usually a human user) and the system itself. A use case thus encapsulates "one specific way of using the system by using some part of the functionality" (p. 154). A set of uses cases and the set of related actors (or user roles) constitute the use case model which is part of the whole requirements model (together with a problem domain object model and user interface descriptions). Use cases thus permit the explicit representation of the intended use of the system and can serve as a basis for discussion with the end users. In this respect they are similar to our technology-enhanced scenarios. However, as use cases represent interactions of single users with the system, the cooperative aspects of using the system are not adequately captured. While it is possible to decompose the cooperative use of a groupware system into several use cases related to different actors, this decomposition omits - especially in the case of synchronous groupware - essential design information like dependencies between different use cases which actually concern the same cooperative system usage. We have made the experience that groupware designers are prone to neglect the cooperative aspects of system use when specifying requirements as use cases. In order to capture the complex dependencies and requirements arising from cooperative work practices, CSCW often draws on methods from the field of Participatory Design (PD). These methods are based on the actual end users taking part in the design process of the system. The first PD projects in Scandinavia were trade union- oriented and focused on giving workers the right to influence their own working conditions, as well as on improving the design of computer systems (For an overview of the history of PD see e.g. the introduction to [8]). However, the obvious advantages of letting those people who know most about the work participate in the design of the systems supporting it, give the methods developed in the context of PD a pragmatic significance beyond their initial political coloring ([16]). PD approaches usually focus on the first part of development processes and spend a lot of time building an understanding of the work context and current work practices of the end users. Sometimes they initially employ ethnographical techniques as a way of gathering socially oriented context information (see e.g. [2]). These techniques rely on project members working very closely with the users in their "natural" surroundings. Not unlike the observation of a tribe's social structures and rules in the Amazon, project members study fields of application like air traffic control centers ([2]) or a control room for the London Underground ([12]). However, ethnographical studies in the proper sense take a lot of time (several month, perhaps years). Thus, their usefulness for commercial (or even short term research) projects has been doubted. [14] discusses different roles which ethnography can play realistically in a design process, ranging from short term evaluation of prior, more long term studies to "quick and dirty" ethnography with the goal of gaining as much insight as possible in a short time. In Holtzblatt and Beyer's Contextual Design methodology [13], the "contextual inquiry" phase ([13], p. 93-94) is an example for the pragmatic use of ethnography in design projects. They suggest that project members - at best the actual designers - spend time (2-3 hours per session) with the users in their usual environment while they are working. The project members can interrupt the work at any time and ask about the goals and motivations behind user actions. In order to provide an efficient view across the whole organization, several project members interview several users in parallel. The interview results are later shared within the project team during a process of structuring what is known about the design problem. Additionally the method employs a variety of work models to capture e.g. the single tasks, steps, or the strategy of the work to be supported. Relevant for CSCW, for instance, is the flow model (p. 97) which depicts the communication between people in the organization. However, the method does not impose a single modeling language on the designer, but proposes to introduce new languages if needed, since there is no one language which can capture all relevant aspects: "Let modeling languages help you. When you must, invent new ones to say exactly what you have to say" (p. 96). Contextual design also involves iteratively going back to the users with clarification questions and design suggestions, e.g. in form of rough paper prototypes (see e.g. [5]) of user interface designs. Even though contextual design does not rely on scenarios in the way the approach presented in this paper does, it is based on a similar understanding of design of interactive systems as a creative and user-centered process. Another approach, which is very much in the tradition of PD, is the MUST method by Kensing et al ([16]). MUST is a Danish acronym for theories of and methods for initial analysis and design activities. MUST specifically supports design in an organizational context and thus explicitly includes steps like strategic analysis which integrates the development efforts into the overall business strategy of the organization. The approach suggest a variety of techniques for developing an understanding of current work practices, including interviews, observations, workshops, document analysis of documents used in the work practice etc. The method also places a strong emphasis on the cooperative development of visions of the future system by both users and developers. Within this process, the use of "scenarios describing envisioned future work practice supported by the proposed design" (p. 137) is suggested. Concerning the use of formal methods, the authors state: ". formalism play a minor role in the MUST method. Instead we suggest plain text, freehand drawings, and sketches for the production and presentation of the relation between proposed IT systems an users' current and future work practice, postponing an extended use of formalism to later on in the development process" (p. 138). Because of its strong organizational focus and its reliance on the PD tradition, the MUST method attempts to explicitly accommodate conflicts of interest between management and workers e.g. in rationalization and downsizing processes. The authors suggest to achieve a consensus concerning the objectives of the system design beforehand. While these issues are out of the scope of this paper, it is nevertheless interesting to note how far the consequences of system design can reach and what responsibilities result for the designers. Scenarios can serve as a tool to make these consequences explicit for everybody involved in the process. A concrete application of scenarios in a CSCW project is discussed by Kyng ([19]), who describes the use of a variety of scenario types in the context of the EUROCOOP project. This project was concerned with designing computer support for cooperation in the Great Belt Link ltd. Company, a state-owned company responsible for the building of a bridge/tunnel between Zealand and Funen in Denmark. It involved the design of four generic, interrelated CSCW applications. Scenarios were used in four different roles: work situation descriptions were supposed to capture relevant, existing situations which the users find to be important parts of their work, bottlenecks, or generally insufficiently supported. These freeform textual descriptions (which Kyng does not yet call scenarios, in contrast to the terminology of this paper) first served as basis for the discussion of current work practices between users and designers. Secondly, they were used in the process of developing mock-up prototypes and accompanying use scenarios which textually describe the intended future use of the envisioned system. In this sense, use scenarios are similar to our technology-enhanced scenarios. In the EUROCOOP project, however, they were mainly employed in setting the stage for user workshops and not for theoretical evaluation. Kyng also describes the project-internal use of more detailed, technically oriented exploration/requirement scenarios. They complemented use scenarios by giving details which are relevant for evaluating technical details (e.g. locking mechanisms) of the proposed design. Finally, explanation scenarios gave a description of the new possibilities offered by the proposed design and the explanation of the rationale behind the design in terms of the working situation. These scenarios were more detailed than use scenarios. Building on the work presented by Kyng, Bardram ([1]) describes the use of scenarios in the SAIK project which was concerned with the design of a Hospital Information System in Denmark. Similar to our work, scenarios were used to describe the current work practices and to envision future, computer supported work activities. Bardram's approach is characterized by the fact that scenarios were used on two different levels of detail. All relevant work practices were textually described as work activity scenarios, while work activities central to system design were documented in greater detail with the help of more structured, tabular analytical scenarios. Furthermore, Bardram points out that current and future scenarios were "alive during the whole system development process" (p. 59) and were constantly updated as the designer's understanding of current work practices and the design itself evolved. In contrast to the role of the use scenarios described by Kyng, analytical scenarios were also employed in a systematic comparison (logical confrontation in Bardram's terminology) of the current work practice and the proposed design. This comparison resulted, for instance, in the discovery of problems concerning the integration of the new system with existing systems for the exchange of EDIFACT messages with external agencies. Additionally, Bardram advocates the use of scenarios in workshops together with the future users and management representatives (compare [16]). This overview of related work shows the span of different roles scenarios can play in the design of systems supporting cooperative work. Scenarios are used to describe current and future, technology supported work practices. They are used for internal, technically-oriented discussions among designers and as a basis for the presentation of prototypes of future systems to end users. They can have different levels of detail and different forms (e.g. tabular or textual). Furthermore, scenarios can be used to support design processes which accommodate organizational and business strategy issues as well as processes which focus on the differences the system makes for individual users (compare our role-based analysis). We conclude that scenarios are an extremely flexible tool for system design and that there is not one single best form and use of scenarios. Depending on the circumstances and constraints of the project, the preferences and experiences of the designers, and last but not least the end users, scenarios can be employed in a number of variations. Summary and Discussion We have presented our use of cooperation scenarios in the design and evaluation of a novel access control system for groupware. Cooperation scenarios are context-rich, informal, textual descriptions of cooperative activities which are gathered and refined through interviews and workshops with user participation. They not only contain a step-by-step description of events but also the goals and subjective opinions (e.g. trust) of persons and other possibly relevant contextual elements. During evaluation we have used scenarios for requirements validation (through communication among designers), role-oriented analysis, and as a basis for realistic workshop presentations. The main advantage of cooperation scenarios lies in the fact that at least some contextual factors can be taken into account early on in the design process. Other, more restrictive forms of scenarios or formal models might omit this qualitative and sometimes rather subjective information. We have shown how cooperation scenarios can be used to identify issues which are important for cooperative activity (trust, awareness, and negotiation). The scenarios allow the designers to think in terms of actual use situations instead of abstract technical criteria. Especially the role-oriented analysis allows a multi-perspective view on a possible design, perhaps helping to identify problems early on (e.g. a benefit and workload disparity). Furthermore, we found that realistic workshop presentations make it possible to establish close contact with the users and get them involved in the evaluation and generation of new ideas in the early stages of the project. On the negative side, we were surprised how easily anecdotal evidence becomes severely distorted due to misunderstandings and misinterpretations of issues in the field of application (e.g. the "virtual" / "physical" desktop problem) and how far apart the different designers' ideas about the users and the field of application can be. We have suggested to address this problem by critically reviewing scenarios among designers and validating them in feedback workshops. Furthermore, we point out again that we see the role of cooperation scenarios early on in the design process when creating innovative functionality and envisioning its effects on cooperative work. Our review of the current literature shows that this reflects the experience made with the use of scenarios in other (CSCW) projects. A good system design subsequently of course makes more formal methods necessary to produce a high quality implementation. Future research We still have a lot to learn about how to design CSCW functionality and how to introduce groupware systems into organizations. Experiences from our projects and reports in the current literature (see e.g. [21]) show that tailorability (or adaptability) is an important success factor for such systems because of the diversified and dynamic nature of requirements of cooperative work. Most current design methodologies still aim at producing a "one-size-fits-all" system design. Part of our current work is concerned with the development of methodologies which not only explicitly specify diversity and dynamics but also help us in deriving the necessary degree of flexibility the final system has to exhibit. Acknowledgements We thank the other past and present members of the POLITeam project at the University of Bonn: Andreas Pfeifer, Helge Kahler, Volkmar Pipek, Markus Won, and Volker Wulf. Furthermore, we are grateful to our POLITeam project partners at GMD in St. Augustin, Germany. Furthermore, the comments of the anonymous reviewers were extremely helpful in improving the final version of this article. The POLITeam project is funded by the BMBF (German Ministry of Research and Education) in the context of the PoliKom research program under grant 01 QA 405 / 0. --R "Scenario-Based Design of Cooperative Systems" "Ethnographically-Informed Systems Design for Air Traffic Control" "Awareness and Coordination in Shared Workspaces" "Cardboard Computers: Mocking-it-up or Hands-on the Future," "Groupware - some Issues and experiences," "Supporting Cooperative Awareness with Local Event Mechanisms: The GoupDesk System" Design at Work. "Data Sharing in Group Work" "Groupware and social dynamics: eight challenges for developers," Reengineering the Corporation - A Manifesto for Business Revolution "Collaboration and Control: Crisis management and multimedia technology in London Underground Line Control Rooms," "Making Customer-Centered Design Work for Teams," "Moving Out from the Control Room: Ethnography in System Design" "MUST - a Method for Participatory Design" "POLITeam - Bridging the Gap between Bonn and Berlin for and with the Users" "Work Processes: Scenarios as a Prelimiary Vocabulary," "Creating Contexts for Design," "Protection," "Experiments with Oval: A Radically Tailorable Tool for Cooperative Work," "Situationsbedingte und benutzerorientierte Anpabarkeit von Groupware," "Access Control for Collaborative Environments" "How to Make Software Softer - Designing Tailorable Applications" "" "On Conflicts and Negotiation in Multiuser Application," --TR --CTR Volker Wulf , Helge Kahler , Volkmar Pipek , Stefan Andiel , Torsten Engelskirchen , Matthias Krings , Birgit Lemken , Meik Poschen , Tim Reichling , Jens Rinne , Markus Rittenbruch , Oliver Stiemerling , Bettina Trpel , Markus Won, ProSEC: research group on HCI and CSCW, ACM SIGGROUP Bulletin, v.21 n.2, p.10-12, August 2000 Steven R. Haynes , Sandeep Purao , Amie L. Skattebo, Situating evaluation in scenarios of use, Proceedings of the 2004 ACM conference on Computer supported cooperative work, November 06-10, 2004, Chicago, Illinois, USA Elizabeth S. Guy, "...real, concrete facts about what works...": integrating evaluation and design through patterns, Proceedings of the 2005 international ACM SIGGROUP conference on Supporting group work, November 06-09, 2005, Sanibel Island, Florida, USA

access control;evaluation;design methodology;groupware;CSCW;cooperation scenarios;scenario-based design

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Convergence to Second Order Stationary Points in Inequality Constrained Optimization.

We propose a new algorithm for the nonlinear inequality constrained minimization problem, and prove that it generates a sequence converging to points satisfying the KKT second order necessary conditions for optimality. The algorithm is a line search algorithm using directions of negative curvature and it can be viewed as a nontrivial extension of corresponding known techniques from unconstrained to constrained problems. The main tools employed in the definition and in the analysis of the algorithm are a differentiable exact penalty function and results from the theory of LC1 functions.

Introduction We are concerned with the inequality constrained minimization problem (P) min f(x) are three times continuously differentiable. Our aim is to develope an algorithm that generates sequences converging to points x satisfying, together with a suitable multiplier - 2 IR m , both the KKT first order necessary optimality conditions (1) and the KKT second order necessary optimality conditions z - 0: (2) In the sequel we will call a point x satisfying (1) a first order stationary point (or just stationary point), while a point satisfying both (1) and (2) will be termed second order stationary point. In the unconstrained case the conditions (1) and (2) boil down to respectively. Standard algorithms for unconstrained minimization usually generate sequences converging to first order stationary points. In a landmark paper [21] (see also [23, 22, 16, 20] and references therein for subsequent developments), McCormick showed that, by using directions of negative curvature in an Armijo-type line search procedure, it is possible to guarantee convergence to second order stationary points. From a theoretical point of view, this is a very strong result, since it makes much more likely that the limit points of the sequence generated by the algorithm are local minimizers and not just saddle points. Furthermore, from a practical point of view, the use of negative curvature directions turns out to be very helpful in the minimization of problems with large non-convex regions [16, 20]. Convergence to second order stationary points was later established also for trust-region algorithms [25, 24], and this constitute one of the main reasons for the popularity of this class of methods. Trust-region algorithms have been extended to equality constrained and box constrained problems so as to mantain convergence to second order stationary points [2, 8, 13, 5, 25, 24, 26]; while negative curvature line search algorithms have been proposed for the linearly inequality constrained case [22, 17]. However, as far as we are aware of, no algorithm for the solution of the more complex nonlinearly inequality constrained minimization Problem (P) exists which generates sequences converging to second order stationary points. The main purpose of this paper is to fill this gap by presenting a negative curvature line search algorithm which enjoys this property. The basic idea behind our approach can be easily explained as follows. (a) Reduce the constrained minimization Problem (P) to an equivalent unconstrained minimization problem by using a differentiable exact penalty function. (b) Apply a negative curvature line search algorithm to the minimization of the penalty function. Although appealingly simple we have to tackle some difficulties to make this approach viable. First of all we have to establish a connection between the unconstrained stationary points of the penalty function provided by the unconstrained minimization algorithm and the constrained second order stationary points of Problem (P). Secondly, we must cope with the fact that differentiable exact penalty functions, although once continuously differentiable, are never twice continuously differentiable everywhere, so that we cannot use an off-the-shelf negative curvature algorithm for its minimization. Furthermore, even in points where the second order derivatives exist, their explicit evaluation would require the use of the third order derivatives of the functions f and g, which we are not willing to calculate. To overcome these difficulties we develop a negative curvature algorithm for the unconstrained minimization of the penalty function which is based on the theory of LC 1 functions and on generalized Hessians. We show that by using a suitable approximation to elements of the generalized Hessian of the penalty function we can guarantee that the unconstrained minimization of the penalty function yields an unconstrained stationary point where a matrix which approximates an element of the generalized Hessian is positive semidefinite. This suffices to ensure that the point so found is also a second order stationary point of Problem (P). We believe that the algorithm proposed in this paper is of intereset because, for the first time, we are able to prove convergence to second order stationary points for general inequality constrained problems. We do so by a fairly natural extension of negative curvature algorithms from unconstrained to constrained problems; we note that the computation of the negative curvature direction can be performed in a manner analogous to and at the same cost as in the unconstrained case. We also remark that we never require the complementarity slackness assumption to establish our results. Finally, we think that the use of some non trivial nonsmooth analysis results to analyze the behavior of smooth algorithms is a novel feature that could be fruitfully applied also in other cases. The paper is organized as follows. In the next section we recall some few known facts about LC 1 functions and generalized Hessians and on the asymptotic identification of active constraints. Furthermore, we also introduce the penalty function along with some of its relevant properties. In Section 3 we introduce and analyze the algorithm. In the fourth section we give some hints on the practical realization of the algorithm. Finally, in the last section we outline possible improvements and make some remarks. We finally review the notation used in this paper. The gradient of a function h : indicated by rh, while its Hessian matrix is denoted by r 2 h. If then the matrix rH is given by I is an index set, with I I is the vector obtained by considering the components of H in I. We indicate by k \Delta k the Euclidean norm and the corresponding matrix norm. If S is a subset of IR n , coS denotes its convex hull. If A is a square matrix, - min (A) denotes its smallest eigenvalue. The Lagrangian of Problem (P) is L(x; by L(x; -(x)) the Lagrangian of Problem (P) evaluated in -(x). Analogously, we indicate by r x L(x; -(x)) (r 2 xx L(x; -(x))) the gradient (Hessian) of the Lagrangian with respect to x evaluated in -(x). Background material In this section we review some results on differentiability of functions and on the identification of active constraints. We also recall the definition and some basic facts about a differentiable exact penalty function for Problem (P) and we establish some related new results. 2.1 LC 1 functions. is said to be an LC 1 function on an open set O if - h is continuously differentiable on O, - rh is locally Lipschitz on O functions were first systematically studied in [18], where the definition of generalized Hessian and the theorem reported below where also given. The gradient of h is locally Lipschitz on O; rh is therefore differentiable almost everywhere in O so that its generalized Jacobian in Clarke's sense [3] can be defined. This is precisely the generalized Hessian of h, whose definition is as follows. be an LC 1 function on the open set O and let x belong to O. We define the generalized Hessian of h at x to be the set @ 2 h(x) of matrices defined as differentiable at x k and r 2 h(x k Note that @ 2 h(x) is a nonempty, convex, compact set of symmetric matrices. Further- more, the point-to-set map x 7! @ 2 h(x) is bounded on bounded sets [18]. For LC 1 functions a second-order Taylor-like expansion is possible. This is the main result on LC 1 function we shall need in this paper. Theorem 2.1 Let h : be an LC 1 function on the open set O and let x and y be two points in O such that [x; y] is contained in O. Then 2.2 Identification of active constraints In this section we recall some results on the identification of active constraints at a stationary point - x of the nonlinear program (P). We refer the interested reader to [14] and to references therein for a detailed discussion of this issue. Here we recall only some results in order to stress the fact that, in a neighborhood of a first order stationary point, it is possible, under mild assumptions, to correctly identify those constraints that are active at the solution. First of all we need some terminology. Given a stationary point - x with a corresponding multiplier - -, which we suppose to be unique, we denote by I 0 (-x) the set of active constraints I 0 (-x) := fij g i while I + (-x) denotes the index set of strongly active constraints I Our aim is to construct a rule which is able to assign to every point x an estimate A(x) so that lies in a suitably small neighborhood of the stationary point - x. Usually estimates of this kind are obtained by comparing the values of g i (x) with the value of an estimate of the multiplier -. For example, it can be easily shown that the set I \Phi (x) := where c is a positive constant and - : is a multiplier function (i.e. a continuous function such that -) coincides with the set I 0 (-x) for all x in a sufficiently small neighborhood of a stationary point - x which satisfies the strict complementarity condition (see the next section for an example of multiplier function). If this condition is violated, then only the inclusions I hold [14]. If the stationary point - x does not satisfy strict complementarity, the situation is therefore more complex, and only recently it has been shown that it is nevertheless possible to correctly estimate the set I 0 (-x) [14]. We will not go into details here, we only point out that the identification of the active constraints when strict complementarity does not hold is possible under very mild assumptions in a simple way. The identification rule takes the following form: where ae(x) is a function that can take different forms according to the assumptions made on the stationary point - x. For example, if in a neighborhood of - x both f and g are analytic (an assumption which is met in most of the practical cases), one can define ae(x) as log(r(x)) where With this choice the set A(x) defined by (4) will coincide with I 0 (-x) in a suitable neighborhood of - x. 2.3 Penalty function In this section we consider a differentiable penalty function for Problem (P), we recall some relevant known facts and prove some new results which are related to the differentiability issues dealt with in Section 2.1. In order to define the differentiable penalty function and to guarantee some of the properties that will be needed we make the following two assumptions. Assumption A. For any x 2 IR n , the gradients rg i (x), i 2 I 0 (x), are linearly independent Assumption B. For any x 2 IR n , the following implication holds Assumptions A and B, together with Assumption C, which will be stated in Section 3, are the only assumptions used to establish the results of this paper. These assumptions, or assumptions similar to them, are frequently encountered in the analysis of constrained minimization algorithms. However, we point out that they can be considerably relaxed; this will be discussed in Section 5. We chose to use this particular set of assumptions in order to simplify the analysis and to concentrate on the issues related to the main topic of the paper, i.e. convergence to second order stationary points. We start by defining a multiplier function where M(x) is the m \Theta m matrix defined by: and G(x) := diag(g i (x)). The main property of this function is that it is continuously differentiable (see below) and, if - x is a first order stationary point, then -x) is the corresponding multiplier (which, by Assumption A, is unique). Using this function we can define the following penalty function is the so-called penalty parameter. Theorem 2.2 The following properties hold: (a) For every ffl, the penalty function Z is continuously differentiable on IR n and its gradient is given by '- where e i is the i-th column of the m \Theta m identity matrix and (x) := diag(- i (x)). (b) For every ffl, the function Z is an LC 1 function on IR n . (c) Let - x be a first order stationary point for Problem (P) and let ffl be given. Then there exists a neighborhood\Omega of - x such that, for every x in\Omega , the following overestimate of the generalized Hessian of Z evaluated at x holds where A is matrix for which we can write kKA (x)k - ae(x) for a nonegative continuous function ae such that In particular, the following overstimate holds in - x where xx L(-x; -x)) +r-A (-x)rg A (-x) T Proof. Point (a) is proved in [11]. Point (b) follows from the expression of the gradient given in (a) taking into account the differentiability assumptions The proof of point (c) can be derived from the very definition of generalized Hessian in the following way. Let - x be a stationary point of Problem (P). Consider a point x in a neighborhood\Omega of - x and sequences of points fx k g converging to x with the gradient of Z existing in x k . This will happen either if (a) for no i or if (b) for all i for which g i I \Phi (x)g. By recalling the expression of the gradient given previously we can write '- '- where I \Psi It is now easy to see that, both in case (a) and (b), the Hessian of Z(x; ffl) in x k can be obtained by differentiating this expression and this gives rg I \Phi where K I \Phi rapresents the sum of terms always containing as a factor either I \Phi Taking into account the definition of @ 2 Z(x; ffl) and that, as discussed in the previous section, if\Omega is suitably small I we have that both g I \Phi x. The assertion of point (c) now follows from these facts and the definition of A. The following theorem gives a sufficient condition, in terms of matrices in the overestimate ~ stationary point of Problem (P) to be a second order stationary point. Theorem 2.3 Let - x be a first order stationary point of Problem (P) and let ffl be given. Then, if a matrix H exists in ~ which is positive semidefinite, - x is a second order stationary point of Problem (P). Proof. Let H in ~ positive semidefinite and suppose by contradiction that x does not satisfy the KKT second order necessary conditions (2). Then a vector z exists such that rg I 0 (recall that -x) equals the multiplier associated with - x). On the other hand, by Theorem 2.2 (c) and by Caratheodory theorem, we also have that, for some integer where, for each i, fi i - 0, A. Since, for each i, A i 2 A, we can write, taking into account the definition of H(-x; ffl; A i ) and (9), z T H(-x; ffl; A i from which z immediately follows. But this contradict the assumption that H is positive semidefinite and the proof is complete. This result will turn out to be fundamental to our approach, since our algorithm will converge to first order stationary stationary points where at least one element in ~ is positive semidefinite. In the remining part of this section we consider some technical results about penalty functions that will be used later on and that help illustrate the relation between the function Z and Problem (P). Proposition 2.4 (a) Let ffl ? 0 and x 2 IR n be given. If x is an unconstrained stationary point of Z and then x is a first order stationary point of Problem (P). (b) Conversely, if x is a first order stationary point of Problem (P), then, for every positive ffl, rZ(x; Proof. See [11]. Proposition 2.5 Let D ae IR n be a compact set. Then, there exists an for every x 2 D and for every ffl 2 (0; -ffl], we have Proof. Let D 1 ae IR n be a compact subset such that D ae intD 1 . In the proof of Proposition 14 in [12] it is shown that if - x is a feasible point in intD 1 , and therefore in D, we can find positive ffl( - x), oe(-x) and ae 0 (-x) such that More precisely, (10) derives from formula (24) in [12] and the discussion which follows that formula. Indicate by M the maximum of krg(x)k for x note that, by Assumption A, M ? 0. Then, recalling that krg(x) T rZ(x; ffl)k - krg(x)kkrZ(x; ffl)k, we can easily deduce from (10) that where we set Suppose now that the theorem is not true. Then, sequences fx k g and fffl k g exist, such that and Since recalling the expression of rZ, gives Thus, by Assumption B we have that - x is feasible. But then, we get a contradiction between (12) and (11), and this concludes the proof. Note that Proposition 2.4 and Proposition 2.5 imply that, given a compact set D, if ffl is sufficiently small, then every stationary point of Problem (P) in D is an unconstrained stationary point of Z and, vice versa, every unconstrained stationary point of Z in D is a stationary point of Problem (P). We refer the interested reader to the review paper [9] and references therein for a more detailed discussion of the properties of differentiable penalty functions. 3 Convergence to second order stationary points In this section we consider a line search algorithm for the minimization of Z(x; ffl) which yields second order stationary points of Problem (P). For the sake of clarity we break the exposition in three parts. In Section 3.1 we first consider a line search algorithm (Algorithm M) which converges, for a fixed value of the penalty parameter ffl, to an unconstrained stationary point of the penalty function Z. By Proposition 2.4 we know that if the penalty parameter were sufficiently small, we would have thus obtained a first order stationary point of Problem (P). Therefore, in Section 3.2, we introduce an algorithm (Algorithm SOC) where Algorithm M is embedded in a simple updating scheme for the penalty parameter based on Proposition 2.5. We show that after a finite number of reductions the penalty parameter stays fixed and every limit point of Algorithm SOC is a first order stationary point of Problem (P). Finally, in Section 3.3 we refine the analysis of Algorithm SOC and we show that every limit point is actually a second order stationary point of Problem (P). In order to establish the results of Sections 3.1, 3.2 and 3.3 we assume that the directions used in Algorithm M satisfy certain conditions. In Section 4 we will illustrate possible ways for generating directions which fulfil these conditions. In order to simplify the analisys we shall assume, from now on, that the following assumption is satisfied. Assumption C. The sequence fx k g of points generated by the algorithms considered below is bounded. 3.1 Convergence for fixed ffl to unconstrained stationary points of Z: Algorithm M We first consider a line search algorithm for the unconstrained minimization of the penalty function Z which generates, for a fixed value ffl of the penalty parameter, sequences converging to unconstrained stationary points of the penalty function. In all this section, ffl is understood to be a fixed positive constant. The algorithm generates a sequence fx k g according to the following rule: Algorithm M where and where ff k is compute by the Linesearch procedure below. Linesearch procedure Step 2: If set ff Step 3: Choose ff 2 [oe 1 ff; oe 2 ff], and go to Step 2. We assume that the matrices H k depend on the sequence fx k g and that the directions and the matrices H k are bounded and satisfy the following conditions: Condition 1. The directions s k are such that rZ(x k ; ffl) T s k - 0 and 0: Condition 2. The directions d k are such that rZ(x k ; ffl) T d k - 0 and, together with the matrices H k , they satisfy ae d T k is not positive semidefinite 0: Condition 3. Let fx k g and fu k g be sequences converging to a first order stationary point - x of Problem (P). Then, for every sequence of matrices fQ k g, with is a sequence of numbers converging to 0. Algorithm M resembles classical line search algorithms using negative curvature directions to force convergence to second order stationary points in unconstrained min- imization. The only apparent difference is that we have the exponent t(x k ) defined by while in corresponding unconstrained algorithms we usually have t(x k every k. We need this change in order to be able to tackle the fact that the penalty function is not everywhere twice continuously differentiable (see, for example, the proof of Proposition 3.1). We also assume that the directions s k , d k and the matrices H k satisfy Conditions 1-3. Conditions 1 and 2 are fairly standard and similar to those employed in the unconstrained case. Condition 3, on the sequence of matrices H k , is, again, related to the nondifferentiability of the gradient of Z. In fact, the matrix H k is supposed to convey some second order information on the penalty function; therefore Condition 3 imposes a certain relation between the matrices H k and the generalized Hessians of Z. Note that if the function Z were twice continuously differentiable the choice would satisfy, by continuity, Condition 3. The following proposition shows that the linesearch procedure described above is well defined in all the cases that, we shall see, are of interest for us. Proposition 3.1 The linesearch procedure is well defined, namely at each iteration the test of Step 2 is satisfied for every ff sufficiently small if the point x k either (a) is not an unconstrained stationary point of the function Z or (b) is a first order stationary point of Problem (P) and H k is not positive semidefinite. Proof. Assume by contradiction the assertion of the proposition is false. Then there exists a sequence fff j g such that ff Z and either the condition (a) or the condition (b) hold. By Theorem 2.1 and taking into account that rZ(x k ; ffl) T d k - 0 by Condition 2, we can find a point Z where Q(u k ) is a symmetric matrix belonging to @ 2 Z(u k ; ffl). Therefore, by (14) and (15), we have: Now we consider two cases. If condition (a) holds, then krZ(x k ; ffl)k 6= 0. We have that dividing both sides of (16) by ff 2 , by taking into account that by making the limit for and by recalling that the sequence fQ(u k )g is bounded, we obtain the contradiction If condition (b) holds, then x k is a first order stationary point of Problem (P) and H k is not positive semidefinite. We have, by Proposition 2.4 (b), that rZ(x k ; so that rZ(x k ; ffl) T s dividing both sides of (16) by ff 2t(x k ) by making the limit for recalling that the sequence fQ(u k )g is bounded, and by recalling that Condition 3 implies with we obtain from (16) which, recalling that H k is not positive semidefinite, contradicts Condition 2. Proposition 3.1 shows that Algorithm M can possibly fail to produce a new point only if, for some k, rZ(x k ; supposing that this trivial case does not occur, the next theorem illustrates the behaviour of an infinite sequence generated by Algorithm M. Theorem 3.2 Let fx k g be an infinite sequence produced by Algorithm M. Then, every limit point x of fx k g is such that rZ(x Proof. Since the sequence fZ(x k ; ffl)g is monotonically decreasing, Z is continuous and fx k g is bounded by Assumption C, it follows that fZ(x k ; ffl)g converges. Hence lim Then, by recalling the acceptability criterion of the line search, Condition 1 and Condition 2, we have: Therefore, (17), (18), Condition 1 and Condition 2 yield: The boundness of s k and d k , Condition 1, Condition 2, (19) and (20) imply in turn: Suppose now, by contradiction, that there exists a converging subsequence fx k gK 1 whose limit point x is not a stationary point. For semplicity and without loss of generality we can rename the subsequence fx k gK 1 by fx k g. Condition 1, (19) and rZ(x ; ffl) 6= 0 imply By (23) we have that there exists an index - k such that, for all k - Z for some oe k 2 [oe Theorem 2.1 and taking into account that rZ(x k ; ffl) T d k - 0 by Condition 2, we can find, for any k - k, a point Z with From (24) and (25) It follows that Dividing both sides by and by simple manipulations we obtain By (21) and (22) we have Condition 1, and since the sequence fQ(u k )g is bounded, we have by (27) lim Condition 1 now implies that rZ(x ; which contradicts the fact the subsequence does not converge to an unconstrained stationary point and this proves the theorem In the next sections, given x k , we indicate by M(x k ) the new point produced by the Algorithm M described above. 3.2 Updating ffl to guarantee convergence to stationary points of Problem (P): Algorithm SOC In this section we show that it is possible to update in a simple way the value of the penalty parameter ffl while minimizing the penalty function Z by Algorithm M, so that every limit point of the sequence of points generated is a first order stationary point of Problem (P). This is accomplished by the Algorithm SOC below. In the next section we shall show that actually, under some additional conditions, the limit points generated by Algorithm SOC are also second order stationary points of Problem (P). This motivates the name SOC, which stands for Second Order Convergence. Algorithm SOC Step 0: Select x 0 and ffl 0 . Set to Step 2; else go to Step 3. Step 2: If max[g(x k and H k 6- 0 go to Step 4; otherwise if max[g(x k Step 3: If krZ(x k go to Step 4; else go to Step 5. Step 4: Compute to Step 1. Step 5: Set x and go to Step 1. Algorithm SOC is related to similar schemes already proposed in the literature (see, e.g., the review paper [9] and references therein). The core step is Step 3, where, at each iteration, the decision of whether to update ffl is taken. This Step is obviously motivated by Propositions 2.5 and Proposition 2.4 (a). Theorem 3.3 Algorithm SOC is well defined. Furthermore, let fx k g and fffl k g be the sequences produced by Algorithm SOC. Then, either the algorithm terminates after p iterations in a first order stationary point x p of Problem (P), or there exist an index - and an -ffl ? 0 such that, for every k - k, ffl and every limit point of the sequence is a first order stationary point of Problem (P). Proof. The algorithm is well defined because every time we reach Step 4 Proposition 3.1 ensures the M(x k ) is well defined. If the algorithm stops after a finite number of iterations, then, by the instructions of Steps 1 and 2, we have rZ(x The thesis then follows by Proposition 2.4 (a). Therefore, assume that an infinite sequence of points is generated. Assumption C and Theorem 2.5 guarantee that ffl is updated only a finite number of times. So, after a finite number of times ffl Algorithm SOC reduces to the application of Algorithm M to Z(x; -ffl). Then, by Theorem 3.2, every limit point - x of fx k g is such that rZ(-x; Since the test at Step 3 is eventually always satisfied, this implies, in turn, that The thesis now follows by Theorem 2.4 (a). 3.3 Algorithm SOC: Second order convergence In this section we prove that under additional suitable conditions, every limit point of the sequence fx k g generated by Algorithm SOC actually satisfies the KKT second order necessary conditions. To establish this result we need the two further conditions below. Condition 4. Let fx k g be a sequence converging to a first order stationary point of Problem (P). Then the directions d k and the matrices H k satisfy Condition 5. Let fx k g be a sequence converging to a first order stationary point - x of Problem (P), and let ffl ? 0 be given. Then Condition 4 mimics similar standard conditions in the unconstrained case, where H k is the Hessian of the objective function. Roughly speaking, it requires the direction d k to be a sufficiently good approximation to an eigenvector corresponding to the smallest eigenvalue of H k . Condition 5, similarly to Condition 3, imposes a connection between the matrices H k and the generalized Hessian of Z. The following theorem establishes the main result of this paper. Theorem 3.4 Let fx k g be the sequence produced by Algorithm SOC. Then, either the algorithm terminates at a second order stationary point x p of Problem (P) or it produces an infinite sequence fx k g such that every limit point x of fx k g is a second order stationary point of Problem (P). Proof. If Algorithm SOC terminates after a finite number of iterations we have, by Theorem 3.3, that x p is a first order stationary point of Problem (P). On the other hand, by the instructions of Step 2 and by Condition 5, we have that H p is positive semidefinite and belongs to ~ Therefore, the assertion follows from Theorem 2.3. We then pass to the case in which an infinite sequence is generated. We already know, by Theorem 3.3, that every limit point of the sequence is a first order stationary point of Problem (P). We also know that eventually ffl k is not updated, so that ffl Then, by Theorem 2.3 it will suffice to show that ~ contains a positive semidefinite element. Suppose the contrary. Let fx k g converge to x . Reasoning as in the beginning of the proof of Theorem 3.2, we have that (21) and still hold. Then, we can assume, renumbering if necessary, that 0: (30) In fact, if this is not the case, (22), Conditions 4 and 5 imply the contradiction that tends to a positive semidefinite element in ~ Then, by (30) and by repeating again the arguments used in the proof of Theorem 3.2, we have that there exists an index - k such that, for all k - k, (26) holds. From (26) we get, recalling Condition 3: which, taking into account that, by Condition 1, rZ(x and the fact that dividing both sides by we have: By (21) and (22) we have so that fQ(u k )g is bounded, while by Condition 5 we have, renumbering if necessary, H by Condition 2, (31) implies : lim and hence, by recalling Condition 4, we have that - min (H Condition 5, contradicts the fact the subsequence fx k g converges to a KKT point where every element in ~ -ffl) is not positive semidefinite. 4 Practical realization In this section we show how we can calculate directions s k , d k and matrices H k satisfying Conditions 1-5 required in the previous sections. Let the matrix H k be defined as and A(x) is any estimate of the active set with the property that, in a neighborhood of a stationary point - x, In section 2.3 we discussed more in detail some possible choices for A(x) and gave adequate references. Note also that, in a stationary point - x, the matrix H k belongs to ~ Given this matrix we have a wide range of choices for s k and d k . theoretically sound option is to take s k to be \GammarZ(x k (b) A more practical choice, however, could be that of taking s k as the solution of the linear system where D k is a diagonal matrix chosen so that the H k +D k is positive definite and the smallest eigenvalue of the sequence fH k +D k g is bounded away from 0. The matrix D k should be 0 if the matrix H k is positive definite with smallest eigenvalue greator than a (small) positive threshold value. Methods for automatically constructing the matrix D k while solving the system H k s are well known and used in the unconstrained case. (c) Another possible choice is to take s k as the direction employed in [10, 1, 15]. can be chosen to be to be an eigenvector associated to the smallest eigenvalue of the matrix H k with the sign possibly changed, in order to ensure rZ(x k ; ffl) T d k - (b) Suitable approximations of the direction of point (a) calculated as indicated, for example, in [23] and [20] could also be employed. The design of an algorithmically effective choice for s k and d k is beyond the scope of this paper. Here we only wanted to illustrate the a wide range of options is available; further choices are certainly possible. In the sequel, for the sake of concreteness, we shall assume that both s k and d k are chosen according to the options (a) listed above. With these choices, and since we are supposing that fx k g remains in a bounded set, it is easy to see that also the sequences are bounded. It is also standard to show that Conditions 1, 2 and 4 are satisfied. Furthermore, if we recall that, in a neighborhood of a stationary point - x of Problem (P), is easy to see that, by the very definition of ~ also Condition 5 is met by our choice for the matrix H k . In the next proposition we show that also the more cumbersome Condition 3 is satisfied. Proposition 4.1 The sequence of matrices defined by (32) satisfies Condition 3. Proof. Let sequences fx k g and fu k g converging to a stationary point of Problem (P) be given. Let fQ k g be any sequence such that Q k 2 @ 2 Z(u k ; ffl) for every k. By Theorem 2.2 (c) we know that we can assume, without loss of generality, that eventually, for x k sufficiently close to the point - x, the matrix Q k has the following form, for some integer where fae k g is a sequence converging to 0 and where, for each i and for each k, fi k A. Since, for each i, A k sufficiently large. We also recall that if A and B are two s \Theta r matrices we can write: r a where a j and b j are the j-th columns of A and B respectively. By employing Taylor expansion we can write r- A k rg A k r- I 0 rg I 0 r- A k rg A k r- I 0 rg I 0 fflB @ r- D k . Now, if we take into account the previous formula and we set a k we can write From this relation the thesis of the proposition readily follows by setting 5 Remarks and conclusions We have presented a negative curvature line search algorithm for the minimization of a nonlinear function subject to nonlinear inequality constraints. The main novel feature of this method is that every limit point of the sequence it generates satisfies both the KKT first and second order necessary optimality conditions. The main tools employed to obtain this result are a continuously differentiable penalty function and some results from the theory of LC 1 functions. For sake of simplicity we did not include equality constraints in our analysis, but they can be easily handled. All the results of this paper go through if one considers also equality constraints, it is sufficient to use an analogous of the penalty function Z where equality constraints are included, see [12]. Another point which deserves attention are the Assumtions A, B and C that we employ. These assumptions are mainly dictated by the penalty function considered; however they can be relaxed if a more sophisticated choice is made for the penalty function. We chose to use the (relatively) simple function Z to concentrate on the main issues related to the second order convergence; however, if the continuously differentiable function proposed in [7] is employed instead of Z, we can improve on the assumptions A, B and C. For example, Assumption A can be relaxed to: For any feasible x, the gradients rg i (x), i 2 I 0 (x) are linearly independent. More significantly, also Assumptions B and C can be considerably relaxed, but to illustrate this point we should introduce some technical notation and we prefere to omit this here and to refer the reader to [7] for more details. We only point out that Assumption C can be replaced by natural and mild assumptions on the problem data which guarantee that the levels sets of the penalty function are compact. --R Constrained Optimization and Lagrange Multiplier Methods. A trust region algorithm for nonlinearly constrained optimization. Optimization and Nonsmooth Analysis. An interior trust region approach for nonlinear minimization subject to bounds. A new trust-region algorithm for equality constrained optimization Global convergence of a class of trust region algorithms for optimization with simple bounds. A continuously differentiable exact penalty function for nonlinear programming problems with unbounded feasible set. On the convergence theory of trust-region- based algorithms for equality-constrained optimization "Algorithms for continuous optimization" "System Modelling and Optimization" A continuously differentiable exact penalty function for nonlinear programming problems with inequalty constraints. Exact penalty functions in constrained optimiza- tion Convergence to a second-order point for a trust-region algorithm with a nonmonotonic penalty parameter for constrained optimization "La Sapienza" Globally and quadratically convergent exact penalty based methods for inequality constrained problems. Nonmonotone curvilinear line search methods for unconstrained optimization. Newton methods for large-scale linear inequality constrained minimization Generalized Hessian matrix and second-order optimality conditions for problems with C 1 New results on a continuously differentiable exact penalty function. "La Sapienza" A modification of Armijo's step-size rule for negative curva- ture Nonlinear Programming: Theory Newton's method with a model trust region modification. --TR --CTR Giovanni Fasano , Massimo Roma, Iterative computation of negative curvature directions in large scale optimization, Computational Optimization and Applications, v.38 n.1, p.81-104, September 2007 Immanuel M. Bomze , Laura Palagi, Quartic Formulation of Standard Quadratic Optimization Problems, Journal of Global Optimization, v.32 n.2, p.181-205, June 2005 X. Q. Yang , X. X. Huang, Partially Strictly Monotone and Nonlinear Penalty Functions for Constrained Mathematical Programs, Computational Optimization and Applications, v.25 n.1-3, p.293-311

inequality constrained optimization;penalty function;KKT second order necessary conditions;LC 1 function;negative curvature direction

297706

A Trace Cache Microarchitecture and Evaluation.

AbstractAs the instruction issue width of superscalar processors increases, instruction fetch bandwidth requirements will also increase. It will eventually become necessary to fetch multiple basic blocks per clock cycle. Conventional instruction caches hinder this effort because long instruction sequences are not always in contiguous cache locations. Trace caches overcome this limitation by caching traces of the dynamic instruction stream, so instructions that are otherwise noncontiguous appear contiguous. In this paper, we present and evaluate a microarchitecture incorporating a trace cache. The microarchitecture provides high instruction fetch bandwidth with low latency by explicitly sequencing through the program at the higher level of traces, both in terms of 1) control flow prediction and 2) instruction supply. For the SPEC95 integer benchmarks, trace-level sequencing improves performance from 15 percent to 35 percent over an otherwise equally sophisticated, but contiguous, multiple-block fetch mechanism. Most of this performance improvement is due to the trace cache. However, for one benchmark whose performance is limited by branch mispredictions, the performance gain is almost entirely due to improved prediction accuracy.

Introduction High performance superscalar processor organizations divide naturally into an instruction fetch mechanism and an instruction execution mechanism. These two mechanisms are separated by instruction issue buffers, for example, issue queues or reservation stations. Conceptually, the instruction fetch mechanism acts as a "producer" which fetches, decodes, and dispatches instructions into the buffer. The instruction execution engine is the "consumer" which issues instructions from the buffer and executes them, subject to data dependence and resource constraints. The instruction issue buffers are collectively called the instruction window. The window is the mechanism for exposing instruction-level parallelism (ILP) in sequential pro- grams: a larger window increases the opportunity for finding data-independent instructions that may issue and execute in parallel. Thus, the trend in superscalar design is to construct larger instruction windows, and provide wider is- sue/execution paths to exploit the corresponding increase in available ILP. These trends place increased demand on the instruction supply mechanism. In particular, the peak instruction fetch rate should match the peak instruction issue rate, or the benefit of aggressive ILP techniques are diminished. In this paper, we are concerned with instruction fetch bandwidth becoming a performance bottleneck. Current fetch units are limited to one branch prediction per cycle and can therefore fetch no more than one basic block per cycle. Previous studies have shown, however, that the average size of basic blocks in integer codes is small, around four to six instructions [30, 3]. While fetching a single basic block each cycle is sufficient for implementations that issue at most four instructions per cycle, it is not for processors with higher peak issue rates. If multiple branch prediction [30, 3, 4, 26] is used, then the fetch unit can at least fetch multiple contiguous basic blocks in a cycle. As will be shown in this paper, fetching multiple contiguous basic blocks is important, but the upper bound on fetch band-width is still limited due to the frequency of taken branches. Therefore, if a taken branch is encountered, it is necessary to fetch instructions down the taken path in the same cycle that the branch is fetched. 1.1. The trace cache The job of the fetch unit is to feed the dynamic instruction stream to the decoder. A problem is that instructions are placed in the cache in their compiled order. Storing programs in this static form favors fetching code with infrequent taken branches or with large basic blocks. Neither of these cases is typical of integer programs. Figure 1(a) shows an example dynamic sequence of basic blocks as they are stored in the instruction cache. The arrows indicate taken branches. Even with multiple branch predictions per cycle, four cycles are required to fetch the instructions in basic blocks ABCDE because the instructions are stored in noncontiguous cache locations. (a) Instruction cache. (b) Trace cache. Figure 1. Storing a noncontiguous sequence of instructions It is for this reason that several researchers have proposed a special instruction cache for capturing long dynamic instruction sequences [15, 22, 23, 24, 21]. This structure is called a trace cache because each line stores a snap- shot, or trace, of the dynamic instruction stream. Referring again to Figure 1, the same dynamic sequence of blocks that appear noncontiguous in the instruction cache are contiguous in the trace cache (Figure 1(b)). The primary constraint on a trace is a maximum length, determined by the trace cache line size. There may be any number of other implementation-dependent constraints, such as the number and type of embedded control transfer instructions, or special terminating conditions for tuning various performance factors [25]. A trace is fully specified by a starting address and a sequence of branch outcomes which describe the path fol- lowed. The first time a trace is encountered, it is allocated a line in the trace cache. The line is filled as instructions are fetched from the instruction cache. If the same trace is encountered again in the course of executing the program, i.e. the same starting address and predicted branch outcomes, it will be available in the trace cache and is fed directly to the decoder in a single cycle. Otherwise, fetching proceeds normally from the instruction cache. Other high bandwidth fetch mechanisms have been proposed that are based on the conventional instruction cache [30, 4, 3, 26]. Every cycle, instructions from noncontiguous locations are fetched from the instruction cache and assembled into the predicted dynamic sequence. This typically requires multiple pipeline stages: (1) a level of indirection through special branch target tables to generate pointers to all of the noncontiguous instruction blocks, (2) a moderate to highly interleaved instruction cache to provide simultaneous access to multiple lines, with the possibility for bank conflicts, and (3) a complex alignment network to shift and align blocks into dynamic program order, ready for decod- ing/renaming. The trace cache approach avoids this complexity by caching dynamic instruction sequences themselves, rather than information for constructing them. If the predicted dynamic sequence exists in the trace cache, it does not have to be recreated on the fly from the instruction cache's static representation. The cost of this approach is redundant instruction storage: the same instructions may reside in both the primary cache and the trace cache, and there is redundancy among different lines in the trace cache. 1.2. Related prior work Alternative High Bandwidth Fetch Mechanisms Four previous studies have focused on mechanisms to fetch multiple, possibly noncontiguous basic blocks each cycle from the instruction cache. These are the branch address cache [30], the subgraph predictor [4], the collapsing buffer [3], and the multiple-block ahead predictor [26]. Trace Cache Development Melvin, Shebanow, and Patt proposed the fill unit and multinodeword cache [18, 16]. The first work qualitatively describes the performance implications of smaller or larger atomic units of work at the instruction-set architecture (ISA), compiler, and hardware levels. The authors argue for small compiler atomic units and large execution atomic units to achieve highest performance. The fill unit is proposed as the hardware mechanism for compacting the smaller compiler units into the large execution units, which are then stored for reuse in a decoded instruction cache. The evaluates the performance potential of large execution atomic units. Although this work only evaluates sizes up to that of a single VAX instruction and a basic block, it also suggests joining two consecutive basic blocks if the intervening branch is "highly predictable". In [17], software basic block enlargement is discussed. In the spirit of trace scheduling [5] and trace selection [11], the compiler uses profiling to identify candidate basic blocks for merging into a single execution atomic unit. The hardware sequences at the level of execution atomic units as created by the compiler. The advantage of this approach is the compiler can optimize and schedule across basic block boundaries. Franklin and Smotherman [6] extended the fill unit's role to dynamically assemble VLIW-like instruction words from a RISC instruction stream, which are then stored in a shadow cache. This structure eases the issue complexity of a wide issue processor. They further applied the fill unit and a decoded instruction cache to improve the decoding performance of a complex instruction-set computer (CISC) [27]. In both cases the cache lines are augmented to store trees to improve the utilization of each line. Four works have independently proposed the trace cache as a complexity-effective approach to high bandwidth instruction fetching. Johnson [15] proposed the expansion cache, which addresses cache alignment, branch prediction throughput, and instruction run merging. The expansion process also predetermines the execution schedule of instructions in a line. Unlike a pure VLIW cache, the schedule may consist of multiple cycles via cycle tagging. Peleg and Weiser [22] describe the design of a dynamic flow instruction cache which stores instructions independent of their virtual addresses, the defining characteristic of trace caches. Rotenberg, Bennett, and Smith [23, 24] motivate the concept with comparisons to other high bandwidth fetch mechanisms proposed in the literature, and defines some of the trace cache design space. Patel, Friendly, and Patt [21] expand upon and present detailed evaluations of this design space, arguing for a more prominent role of the trace cache. The mispredict recovery cache proposed by Bondi, Nanda, and Dutta [1] caches instruction threads from alternate paths of mispredicted branches. The goal of this work is to quickly bypass the multiple fetch and decode stages of a long CISC pipeline following a branch mispredict. Nair and Hopkins [19] employ dynamic instruction formatting to cache large scheduled groups, similar in spirit to the cycle tagging approach of the expansion cache. There has also been recent work incorporating trace caches into new processing models. Vajapeyam and Mitra [29], Sundararaman and Franklin [28], and Rotenberg, Jacobson, Sazeides, and Smith [25] exploit the data and control hierarchy implied by traces to overcome complexity and architectural hurdles of superscalar processors. Jacob- son, Rotenberg, and Smith [14] propose a control prediction model well suited to the trace cache called next trace pre- diction, discussed in later sections. Friendly, Patel, and Patt propose a new processing model called inactive issue for reducing the effects of branch mispredictions [7], and dynamically optimizing traces before storing them in the trace cache, reducing their execution time significantly [8]. Microcode, VLIW, and Block-Structured ISAs Clearly the concept of traces exists in the software realm of instruction-level parallelism. Early work by Fisher [5], Hwu and Chang [11], and others on trace scheduling and trace selection for microcode recognized the problem imposed by branches on code optimization. Subsequent VLIW architectures and novel ISA techniques, for example [12, 10], further promote the ability to schedule long sequences of instructions containing multiple branches. 2. Trace cache microarchitecture In Section 1.1 we introduced the concept of the trace cache - an instruction cache which captures dynamic instruction sequences, or traces. We now present a microarchitecture organized around traces. 2.1. Trace-level sequencing The premise of the proposed microarchitecture, shown in Figure 2, is to provide high instruction fetch bandwidth with low latency. This is achieved by explicitly sequencing through the program at the higher level of traces, both for (1) control flow prediction and (2) supplying instructions. Cache Instruction Branch Trace Trace Cache outstanding trace buffers Execution Engine branch outcomes update Figure 2. Microarchitecture. A next trace predictor [14] treats traces as basic units and explicitly predicts sequences of traces. Because traces are the unit of prediction, rather than individual branches, high branch prediction throughput is implicitly achieved with only a single trace prediction per cycle. Jacobson et al [14] demonstrated that explicit trace prediction not only removes fundamental constraints on the number of branches in a trace (usually a consequence of adapting single branch predictors to multiple branch predictor counterparts [23]), but it also holds the potential for achieving higher overall branch prediction accuracy than single branch predictors. Details of next trace prediction are presented in Section 2.3. The output of the trace predictor is a trace identifier: a given trace is uniquely identified by its starting PC and the outcomes of all conditional branches embedded in the trace. The trace identifier is used to lookup the trace in the trace cache. The index into the trace cache can be derived from just the starting PC, or a combination of PC and branch out- comes. Using branch outcomes in the index has the advantage of providing path associativity - multiple traces emi- nating from the same start PC can reside simultaneously in the trace cache even if it is direct mapped [24]. The output of the trace cache is one or more traces, depending on the cache associativity. A trace identifier is stored with each trace in order to determine a trace cache hit, analogous to the tag of conventional caches. The desired trace is present in the cache if one of the cached trace identifiers matches the predicted trace identifier. The trace predictor and trace cache together provide fast trace-level sequencing. Unfortunately, trace-level sequencing does not always provide the required trace. This is particularly true at the start of the program or when a new region of code is reached - neither the trace predictor nor the trace cache has "learned" any traces yet. Instruction-level sequencing, discussed in the next section, is required to construct non-existent traces or repair trace mispredictions. 2.2. Instruction-level sequencing The outstanding trace buffers in Figure 2 are used to (1) construct new traces that are not in the trace cache and (2) track branch outcomes as they become available from the execution engine, allowing detection of mispredictions and repair of the traces containing them. Each fetched trace is dispatched to both the execution engine and an outstanding trace buffer. In the case of a trace cache miss, only the trace prediction is received by the allocated buffer. The trace prediction itself provides enough information to construct the trace from the instruction cache, although this typically requires multiple cycles due to predicted-taken branches. In the case of a trace cache hit, the trace is dispatched to the buffer. This allows repair of a partially mispredicted trace, i.e. when a branch outcome returned from execution does not match the path indicated within the trace. In the event of a branch misprediction, the trace buffer begins reconstructing the tail of the trace (or all of the trace if the start PC is incorrect) using the corrected branch target and the instruction cache. For subsequent branches in the trace, a second-level branch predictor is used to make predictions. We advocate an aggressive instruction cache design for providing robust performance over a broad range of trace cache miss rates. The instruction cache is 2-way interleaved so that up to a full cache line can be fetched each cycle, independent of PC alignment [9]. The second-level branch prediction mechanism is simple - a 2-bit counter and branch target stored with each branch. Logically, the instructions, counters, and targets are all stored in the instruction cache (as opposed to a separate cache and branch target buffer) to allow fast, parallel prediction of any number of not-taken branches. We call this instruction fetch mechanism SEQ.n in keeping with the terminology of [24] - any number (de- noted n) of sequential basic blocks, up to the line size, can be fetched in a single cycle. When a trace buffer is through constructing its trace, it is written into the trace cache and dispatched to the execution engine. If the newly constructed trace is a result of misprediction recovery, the trace identifier is also sent to the trace predictor for repairing its path history. 2.3. Next trace prediction The next trace predictor, shown in Figure 3, is based on Jacobson's work on path-based, high-level control flow prediction [13, 14]. An index into a correlated prediction table is formed from the sequence of past trace identifiers. The hash function used to generate the index is called a DOLC func- tion: 'D'epth specifies the path history depth in terms of traces; 'O'ldest indicates the number of bits selected from each trace identifier except the two most recent ones; 'L'ast path PC path PC path HASH PC, path Prediction Table path to Trace Cache Figure 3. Jacobson's next trace predictor. and 'C'urrent indicate the number of bits selected from the second-most recent and most recent trace identifiers, respectively Each entry in the correlated prediction table contains a trace identifier and a 2-bit counter for replacement. The predictor is augmented with several other mechanisms [14]. ffl Hybrid prediction. In addition to the correlated table, a second, smaller table is indexed with only the most recent trace identifier. This second table requires a shorter learning time and suffers less aliasing pressure. ffl Return history stack. At call instructions, the path history is pushed onto a special stack. When the corresponding return point is reached, path history before the call is restored. This improves accuracy because control flow following a subroutine is highly correlated with control flow before the call. Alternate trace identifier. An entry in the correlated table may be augmented with an alternate trace predic- tion, a form of associativity in the predictor. If a trace misprediction is detected, the outstanding trace buffer responsible for repairing the trace can use the alternate prediction if it is consistent with known branch outcomes in the trace. If so, the trace buffer does not have to resort to the second-level branch predic- instruction-level sequencing is avoided altogether if the alternate trace also hits in the trace cache. 2.4. Trace selection The performance of the trace cache is strongly dependent on trace selection, the algorithm used to divide the dynamic instruction stream into traces. Trace selection primarily affects average trace length and trace cache hit rate, both of which, in turn, affect fetch bandwidth. The interaction between trace length and hit rate, however, is not well under- stood. Preliminary studies indicate that longer traces result in lower hit rates, but this may be an artifact of naive trace selection policies. Sophisticated selection techniques that are conscious of control flow constructs - loop back-edges, loop fall-through points, call sites, and re-convergent points in general - may lead to different conclusions. The reader is referred to [21, 25, 20] for a few interesting control-flow- conscious selection heuristics. Trace selection in this paper is constrained only by the maximum trace length of 16 instructions, and indirect branches (returns and jump/call indirects) terminate traces. 2.5. Hierarchical sequencing In Figure 4(a), a portion of the dynamic instruction stream is shown with a solid horizontal arrow from left to right. The stream is divided into traces through T5. This sequence of traces is produced independent of where the instructions come from - trace predictor/trace cache, trace predictor/instruction cache, or branch predictor/instruction cache. mispredicted branch (a) Hierarchical. mispredicted branch (b) Non-hierarchical. Figure 4. Two sequencing models. For example, if the trace predictor mispredicts T3, the trace buffer assigned to T3 resorts to instruction-level se- quencing. This is shown in the diagram as a series of steps, depicting smaller blocks fetched from the instruction cache. The trace buffer strictly adheres to the boundary between T3 and T4, dictated by trace selection, even if the final instruction cache fetch produces a larger block of sequential instructions than is needed by T3 itself. We call this process hierarchical sequencing because there exists a clear distinction between inter-trace control flow and intra-trace control flow. Inter-trace control flow, i.e. trace boundaries, is effectively pre-determined by trace selection and is unaffected by dynamic effects such as trace cache misses and mispredictions. A contrasting sequencing model is shown in Figure 4(b). In this model, trace selection is "reset" at the point of the mispredicted branch, producing the shifted traces T3 0 , T4 0 , and T5 0 . This sequencing model does not work well with path-based next trace prediction. After resolving the branch misprediction, trace T3 0 and subsequent traces must somehow be predicted. However, this requires a sequence of traces leading to T3 0 and no such sequence is available (in- dicated with question marks in the diagram). A potential problem with hierarchical sequencing is mis-prediction recovery latency. Explicit next trace prediction uses a level of indirection: a trace is first predicted, and then the trace cache is accessed. This implies an extra cycle is added to the latency of misprediction recovery. How- ever, this extra cycle is not exposed. First, consider the case in which the alternate trace prediction is used. The primary and alternate predictions are supplied by the trace predictor at the same time, and stored together in the trace buffer. Therefore, the alternate prediction is immediately available for accessing the trace cache when the misprediction is detected. Second, if the alternate is not used, then the second-level branch predictor and instruction cache are used to fetch instructions from the correct path. In this case, the instruction cache is accessed immediately with the correct branch target PC returned by the execution engine. In our evaluation, we assume a trace must be fully constructed before any of its instructions are dispatched to the execution engine, because traces are efficiently renamed as a unit [29, 25]. This aggravates both trace misprediction and trace cache miss recovery latency. We want to make it clear, however, that this is not due to any fundamental constraint of the fetch model, only an artifact of our dispatch model. 3. Simulation methodology 3.1. Fetch models To evaluate the performance of the trace cache microar- chitecture, we compare it to several more constrained fetch models. We first determine the performance advantage of fetching multiple contiguous basic blocks per cycle over conventional single block fetching. Then, the benefit of fetching multiple noncontiguous basic blocks is isolated. In all models a next trace predictor is used for control prediction, for two reasons. First, next trace prediction is highly accurate, and whether predicting one or many branches at a time, it is comparable to or better than some of the best single branch predictors in the literature. Second, it is desirable to have a common underlying predictor for all fetch models so we can separate performance due to fetch bandwidth from that due to branch prediction (more on this in Section 3.2). What differentiates the following models is the trace selection algorithm. ffl SEQ.1 ("sequential, 1 block"): A "trace" is a single basic block up to 16 instructions in length. ffl SEQ.n ("sequential, n blocks"): A "trace" may contain any number of sequential basic blocks up to the instruction limit. ("trace cache"): A trace may contain any number of conditional branches, both taken and not-taken, up to instructions or the first indirect branch. The SEQ.1 and SEQ.n models do not use a trace cache because an interleaved instruction cache is capable of supplying a "trace" in a single cycle [9] - a consequence of the sequential selection constraint. Therefore, one may view the SEQ.1/SEQ.n fetch unit as identical to the trace cache microarchitecture in Figure 2, except the trace cache block is replaced with a conventional instruction cache. That is, the next trace predictor drives a conventional instruction cache, and the trace buffers are used to construct "traces" from the L2 cache/main memory if not present in the cache. Finally, to establish an upper bound on the performance of noncontiguous instruction fetching, we introduce a fourth model, TC-perfect, which is identical to TC but the trace cache always hits. 3.2. Isolating trace predictor/trace cache performance An interesting side-effect of trace selection is that it significantly affects trace prediction accuracy. In general, smaller traces (resulting from more constrained trace selec- tion) result in lower accuracy. We have determined at least two reasons for this. First, longer traces naturally capture longer path history. This can be compensated for by using more trace identifiers in the path history if the traces are small; that is, a good DOLC function for one trace length is not necessarily good for another. For the TC model, DOLC (a depth of 7 traces) consistently performs well over all benchmarks [14]. For SEQ.1 and SEQ.n, a brief search of the design space shows depth of 17 traces) performs well. We have observed, however, that tuning the DOLC parameters is not enough - trace selection affects accuracy in other ways. The graph in Figure 5 shows trace predictor performance using an unbounded table, i.e. using full, unhashed path history to make predictions. The graph shows trace mispredictions per 1000 instructions for SEQ.1, SEQ.n, and TC trace selection, as the history depth is var- ied. For the go benchmark, trace mispredictions for the SEQ.n model do not dip below 8.8 per 1000 instructions, whereas the TC model reaches as few as 8.0 trace mispredictions per 1000 instructions. Unconstrained trace selection results in the creation of many unique traces. While this trace explosion generally has a negative impact on trace cache performance, we hypothesize it also creates many more unique contexts for making predictions. A large prediction table can exploit this additional context. ideal trace prediction (GO)913171 6 11 history depth (traces) trace misp/1000 instr. SEQ.n Figure 5. Impact of trace selection on unbounded trace predictor performance. We conclude that it is difficult to separate the performance advantage of the trace cache from that of the trace predictor, because both show positive improvement with longer traces. Nonetheless, when we compare TC to SEQ.n or SEQ.1, we would like to know how much benefit is derived from the trace cache itself. To this end, we developed a methodology to statistically "adjust" the overall branch prediction accuracy of a given fetch model to match that of another model. The trace predictor itself is not adjusted - it produces predictions in the normal fashion. However, after making a prediction, the predicted trace is compared with the actual trace, determined in advance by a functional simulator running in parallel with the timing simulator. If the prediction is in- correct, the actual trace is substituted for the mispredicted trace with some probability. In other words, some fraction of mispredicted traces are corrected. The probability for injecting corrections was chosen on a per-benchmark basis to achieve the desired branch misprediction rate. This methodology introduces two additional fetch mod- els, SEQ.1-adj and SEQ.n-adj, corresponding to the "ad- justed" SEQ.1 and SEQ.n models. Clearly these models are unrealizable, but they are useful for performance comparisons because their adjusted branch misprediction rates match that of the TC model. 3.3. Simulator and benchmarks A detailed, fully-execution driven superscalar processor simulator is used to evaluate the trace cache microarchitec- ture. The simulator was developed using the simplescalar platform [2]. This platform uses a MIPS-like instruction set and a gcc-based compiler to create binaries. The datapath of the fetch engine as shown in Figure 2 is faithfully modeled. The next trace predictor has 2 tries. The DOLC functions for compressing the path history into a 16-bit index were described earlier in Section 3.2, for both the TC and SEQ models. The trace cache configuration - size, associativity, and indexing - is varied. There are sufficient outstanding trace buffers to keep the instruction window full. The trace buffers share a single port to the combined instruction cache and second-level branch predic- tor. The instruction cache is 64KB, 4-way set-associative, and 2-way interleaved. The line size is 16 instructions and the cache hit and miss latencies are 1 cycle and 12 cycles respectively. The second-level branch predictor consists of 2-bit counters and branch targets, assumed to be logically stored with each branch in the instruction cache. An instruction window of 256 instructions is used in all experiments. The processor is 16-way superscalar, i.e. the processor can fetch and issue up to 16 instructions each cycle. Five basic pipeline stages are modeled. Instruction fetch and dispatch take 1 cycle each. Issue takes at least 1 cycle, possibly more if the instruction must stall for operands; any 16 instructions, including loads and stores, may issue each cycle. Execution takes a fixed latency based on instruction type, plus any time spent waiting for a result bus. Instructions retire in order. For loads and stores, address generation takes 1 cycle and the cache access is 2 cycles for a hit. The data cache is 64KB, 4-way set-associative with a line size of 64 bytes and a miss penalty of 14 cycles. Realistic but aggressive memory disambiguation is modeled. Loads may proceed ahead of any unresolved stores, and any memory hazards are detected as store addresses become available - recovery is via selective reissuing of misspeculated loads and their dependent instructions [25]. Seven of the SPEC95 integer benchmarks, shown in Table are simulated to completion. Table 1. Benchmarks. benchmark input dataset dynamic instr. count gcc -O3 genrecog.i 117M jpeg vigo.ppm 166M li queens 7 202M perl scrabbl.pl < scrabbl.in 108M vortex persons.250 101M 4. Results 4.1. Performance of fetch models Figure 6 shows the performance of the six fetch models in terms of retired instructions per cycle (IPC). The model in this section uses a 64KB (instruction storage only), 4-way set-associative trace cache. The trace cache is indexed using only the PC (i.e. no explicit path associativity, except that afforded by the 4 ways).357go gcc jpeg li perl m88k vortex IPC TC-perfect Figure 6. Performance of the fetch models. We can draw several conclusions from the graph in Figure 6. First, comparing the SEQ.n models to the SEQ.1 models, it is apparent that predicting and fetching multiple sequential basic blocks provides a significant performance advantage over conventional single-block fetching. The graph in Figure 7 shows that the performance advantage of the SEQ.n model over the SEQ.1 model ranges from about 5% to 25%, with the majority of benchmarks showing greater than 15% improvement. Similar results hold whether or not branch prediction accuracy is adjusted for the SEQ.n and SEQ.1 models. This first observation is important because the SEQ.n model only requires a more sophisticated, high-level control flow predictor, and retains a more-or-less conventional instruction cache microarchitecture. 0% 5% 10% 15% 20% 30% go gcc jpeg li perl m88k vort improvement in IPC SEQ.n over SEQ.1 SEQ.n-adj over SEQ.1-adj Figure 7. Speedup of SEQ.n over SEQ.1. Second, the ability to fetch multiple, possibly noncontiguous basic blocks improves performance significantly over sequential-only fetching. The graph in Figure 8 shows that the performance advantage of the TC model over the SEQ.n model ranges from 15% to 35%. Speedup of TC over SEQ.n 0% 5% 10% 15% 20% 30% 40% go gcc jpeg li perl m88k vort improvement in IPC trace cache trace predictor Figure 8. Speedup of TC over SEQ.n. Figure 8 also isolates the contributions of next trace prediction and the trace cache to performance. The lower part of each bar is the speedup of model SEQ.n-adj over SEQ.n. And since the overall branch misprediction rate of SEQ.n- adj is adjusted to match that of the TC model, this part of the bar approximately isolates the impact of next trace prediction on performance. The top part of the bar therefore isolates the impact of the trace cache on performance. For go, which suffers noticeably more branch mispredictions than other benchmarks, most of the benefit of the model comes from next trace prediction. In this case, the longer traces of the TC model are clearly more valuable for improving the context used by the next trace predictor than for providing raw instruction bandwidth. For gcc, however, both next trace prediction and the trace cache contribute equally to performance. The other five benchmarks benefit mostly from higher fetch bandwidth. Finally, Figure 6 shows the moderately large trace cache of the TC model very nearly reaches the performance upper bound established by TC-perfect (within 4%). Table shows trace- and branch-related measures. Average trace lengths for TC range from 12.4 (li) to 15.8 (jpeg) instructions (1.6 to over 2 times longer than SEQ.n traces). The table also shows predictor performance: primary and alternate trace mispredictions per 1000 instructions, and overall branch misprediction rates (the latter is computed by checking each branch at retirement to see if it caused a mis- prediction, whether originating from the trace predictor or second-level branch predictor). In all cases prediction improves with longer traces. TC has from 20% to 45% fewer trace mispredictions than SEQ.1, resulting in 15% (jpeg) to 41% (m88ksim) fewer total branch mispredictions. Note that the adjusted branch misprediction rates for the SEQ models are nearly equal to those of TC. Shorter traces, however, generally result in better alternate trace prediction accuracy. Shorter traces result in (1) fewer total traces and thus less aliasing, and (2) fewer possible alternative traces from a given starting PC. For all benchmarks except gcc and go, the alternate trace prediction is almost always correct given the primary trace prediction is incorrect - both predictions taken together result in fewer than 1 trace misprediction per 1000 instructions. Trace caches introduce redundancy - the same instruction can appear multiple times in one or more traces. Table 2 shows two redundancy measures. The overall redundancy factor, RF overall , is computed by maintaining a table of all unique traces ever retired. Redundancy is the ratio of total number of instructions to total number of unique instructions for traces collected in the table. RF overall is independent of trace cache configuration and does not capture dynamic behavior. The dynamic redundancy factor, RF dyn , is computed similarly, but using only traces in the trace cache in a given cycle; the final value is an average over all cycles. RF dyn was measured using a 64KB, 4-way trace cache. RF overall varies from 2.9 (vortex) to 14 (go). RF dyn is less than RF overall and only ranges between 2 and 4, because the fixed size trace cache limits redundancy, and perhaps temporally there is less redundancy. 4.2. Trace cache size and associativity In this section we measure performance of the TC model as a function of trace cache size and associativity. Figure 9 shows overall performance (IPC) for 12 trace cache config- urations: direct mapped, 2-way, and 4-way associativity for each of four sizes, 16KB, 32KB, 64KB, and 128KB. Associativity has a noticeable impact on performance for Table 2. Trace statistics. model measure gcc go jpeg li m88k perl vort trace length 4.9 6.2 8.3 4.2 4.8 5.1 5.8 trace misp./1000 8.8 14.5 5.2 6.9 3.5 3.4 1.5 SEQ.1 alt. trace misp./1000 2.1 4.5 branch misp. rate 5.0% 11.0% 7.7% 3.7% 2.2% 2.2% 1.1% adjusted misp. rate 3.6% 8.2% 6.6% 3.2% 1.3% 1.4% 0.8% trace length 7.2 8.0 9.6 6.3 6.0 7.1 8.2 trace misp./1000 7.3 12.7 4.6 6.9 3.3 3.1 1.2 SEQ.n alt. trace misp./1000 2.7 5.4 0.5 0.9 0.6 0.3 0.3 branch misp. rate 4.4% 10.1% 7.0% 3.7% 2.1% 2.0% 0.9% adjusted misp. rate 3.6% 8.1% 6.7% 3.1% 1.3% 1.4% 0.8% trace length 13.9 14.8 15.8 12.4 13.1 13.0 14.4 trace misp./1000 5.4 9.6 4.2 5.5 2.0 2.1 1.0 alt. trace misp./1000 2.7 5.3 0.9 1.3 0.5 0.3 0.3 branch misp. rate 3.6% 8.2% 6.7% 3.1% 1.3% 1.5% 0.8% control instr. per trace 2.8 2.3 1.3 2.9 2.5 2.5 2.3 RF overall 7.1 14.4 5.3 3.1 3.7 4.1 2.9 RF dyn 3.0 3.3 3.7 3.2 3.1 2.9 2.13.54.55.56.57.5 trace cache size IPC jpeg perl vort go gcc li4 Figure 9. Performance vs. size/associativity. all of the benchmarks except go. Go has a particularly large working set of unique traces [25], and total capacity is more important than individual trace conflicts. The curves of jpeg and li are fairly flat - size is of little importance, yet increasing associativity improves performance. These two benchmarks suffer few general conflict misses (otherwise size should improve performance), yet conflicts among traces with the same start PC are significant. Associativity allows simultaneously caching these path-associative traces. The performance improvement of the largest configuration (128KB, 4-way) with respect to the smallest one (16KB, direct mapped) ranges from 4% (go) to 10% (gcc). Figure shows trace cache performance in misses per 1000 instructions. Trace cache size is varied along the x- axis, and there are six curves: direct mapped (DM), 2- way (2W), and 4-way (4W) associative caches, both with and without indexing for path associativity (PA). We chose (somewhat arbitrarily) the following index function for achieving path associativity: the low-order bits of the PC form the set index, and then the high-order bits of this index are XORed with the first two branch outcomes of the trace identifier. Gcc and go are the only benchmarks that do not fit entirely within the largest trace cache. As we observed earlier, go has many heavily-referenced traces, resulting in no fewer than 20 misses/1000 instructions. Path associativity reduces misses substantially, particularly for direct mapped caches. Except for vortex, path associativity closes the gap between direct mapped and 2-way associative caches by more than half, and often entirely. misses/1000 instr GO253545 misses/1000 instr LI51525M88KSIM2610misses/1000 instr PERL5152535 trace cache size trace cache size misses/1000 instr DM DM-PA 4W-PA Figure 10. Trace cache misses. 5. Summary It is important to design instruction fetch units capable of fetching past multiple, possibly taken branches each cycle. Trace caches provide this capability without the complexity and latency of equivalent-bandwidth instruction cache designs. We evaluated a microarchitecture incorporating a trace cache, with the following major results. ffl The trace cache improves performance from 15% to 35% over an otherwise equally-sophisticated, but contiguous multiple-block fetch mechanism. ffl Longer traces improve trace prediction accuracy. For the misprediction-bound benchmark go, this factor contributes almost entirely to the observed performance gain. ffl A moderately large and associative trace cache performs as well as a perfect trace cache. For go, however, trace mispredictions mask poor trace cache performance. ffl Overall performance is not as sensitive to trace cache size and associativity as one might expect, due in part to robust instruction-level sequencing. IPC varies no more than 10% over a wide range of configurations. ffl The complexity advantage of the trace cache comes at the price of redundant instruction storage: for gcc, a factor of 7 redundancy among all traces created, corresponding to a factor of 3 redundancy in the trace cache. ffl An instruction cache combined with an aggressive trace predictor can fetch any number of contiguous basic blocks per cycle, yielding from 5% to 25% improvement over single-block fetching. Acknowledgments Our research on trace caches had its genesis in stimulating group discussions with Guri Sohi and his students Todd Austin, Scott Breach, Andreas Moshovos, Dionisios Pnevmatikatos, and T. N. Vijaykumar; their contribution is gratefully acknowledged. We would also like to give special thanks to Quinn Jacobson for his valuable input and for providing access to next trace prediction simulators. This work was supported in part by NSF Grants MIP- 9505853 and MIP-9307830 and by the U.S. Army Intelligence Center and Fort Huachuca under Contract DABT63- 95-C-0127 and ARPA order no. D346. Eric Rotenberg is supported by an IBM Fellowship. --R Integrating a misprediction recovery cache (mrc) into a superscalar pipeline. Evaluating future mi- croprocessors: The simplescalar toolset Optimization of instruction fetch mechanisms for high issue rates. Control flow prediction with tree-like subgraphs for superscalar processors Trace scheduling: A technique for global microcode compaction. Alternative fetch and issue policies for the trace cache fetch mechanism. Putting the fill unit to work: Dynamic optimizations for trace cache microproces- sors Branch and fixed-point instruction execution units Increasing the instruction fetch rate via block-structured instruction set ar- chitectures Trace selection for compiling large c application programs to microcode. Control flow speculation in multiscalar processors. Expansion caches for superscalar processors. Performance benefits of large execution atomic units in dynamically scheduled machines. Exploiting fine-grained parallelism through a combination of hardware and software techniques Hardware support for large atomic units in dynamically scheduled machines. Exploiting instruction level parallelism in processors by caching scheduled groups. Improving trace cache effectiveness with branch promotion and trace packing. Critical issues regarding the trace cache fetch mechanism. Dynamic flow instruction cache memory organized around trace segments independent of virtual address line. Trace cache: a low latency approach to high bandwidth instruction fetch- ing Trace cache: a low latency approach to high bandwidth instruction fetch- ing Trace processors. Improving cisc instruction decoding performance using a fill unit. Multiscalar execution along a single flow of control. Improving superscalar instruction dispatch and issue by exploiting dynamic code se- quences Increasing the instruction fetch rate via multiple branch prediction and a branch address cache. --TR --CTR Emil Talpes , Diana Marculescu, Power reduction through work reuse, Proceedings of the 2001 international symposium on Low power electronics and design, p.340-345, August 2001, Huntington Beach, California, United States S. Bartolini , C. A. Prete, A proposal for input-sensitivity analysis of profile-driven optimizations on embedded applications, ACM SIGARCH Computer Architecture News, v.32 n.3, p.70-77, June 2004 Emil Talpes , Diana Marculescu, Execution cache-based microarchitecture power-efficient superscalar processors, IEEE Transactions on Very Large Scale Integration (VLSI) Systems, v.13 n.1, p.14-26, January 2005 Emil Talpes , Diana Marculescu, Increased Scalability and Power Efficiency by Using Multiple Speed Pipelines, ACM SIGARCH Computer Architecture News, v.33 n.2, p.310-321, May 2005 Michael Behar , Avi Mendelson , Avinoam Kolodny, Trace cache sampling filter, ACM Transactions on Computer Systems (TOCS), v.25 n.1, p.3-es, February 2007 Oliverio J. Santana , Ayose Falcn , Alex Ramirez , Mateo Valero, Branch predictor guided instruction decoding, Proceedings of the 15th international conference on Parallel architectures and compilation techniques, September 16-20, 2006, Seattle, Washington, USA Oliverio J. Santana , Alex Ramirez , Josep L. Larriba-Pey , Mateo Valero, A low-complexity fetch architecture for high-performance superscalar processors, ACM Transactions on Architecture and Code Optimization (TACO), v.1 n.2, p.220-245, June 2004 Sang-Jeong Lee , Pen-Chung Yew, On Augmenting Trace Cache for High-Bandwidth Value Prediction, IEEE Transactions on Computers, v.51 n.9, p.1074-1088, September 2002 Xianglong Huang , Stephen M. Blackburn , David Grove , Kathryn S. McKinley, Fast and efficient partial code reordering: taking advantage of dynamic recompilatior, Proceedings of the 2006 international symposium on Memory management, June 10-11, 2006, Ottawa, Ontario, Canada S. Bartolini , C. A. Prete, Optimizing instruction cache performance of embedded systems, ACM Transactions on Embedded Computing Systems (TECS), v.4 n.4, p.934-965, November 2005 Yoav Almog , Roni Rosner , Naftali Schwartz , Ari Schmorak, Specialized Dynamic Optimizations for High-Performance Energy-Efficient Microarchitecture, Proceedings of the international symposium on Code generation and optimization: feedback-directed and runtime optimization, p.137, March 20-24, 2004, Palo Alto, California Michele Co , Dee A. B. Weikle , Kevin Skadron, Evaluating trace cache energy efficiency, ACM Transactions on Architecture and Code Optimization (TACO), v.3 n.4, p.450-476, December 2006 independence in trace processors, Proceedings of the 32nd annual ACM/IEEE international symposium on Microarchitecture, p.4-15, November 16-18, 1999, Haifa, Israel Roni Rosner , Yoav Almog , Micha Moffie , Naftali Schwartz , Avi Mendelson, Power Awareness through Selective Dynamically Optimized Traces, ACM SIGARCH Computer Architecture News, v.32 n.2, p.162, March 2004 Roni Rosner , Micha Moffie , Yiannakis Sazeides , Ronny Ronen, Selecting long atomic traces for high coverage, Proceedings of the 17th annual international conference on Supercomputing, June 23-26, 2003, San Francisco, CA, USA Alex Ramirez , Oliverio J. Santana , Josep L. Larriba-Pey , Mateo Valero, Fetching instruction streams, Proceedings of the 35th annual ACM/IEEE international symposium on Microarchitecture, November 18-22, 2002, Istanbul, Turkey

multiple branch prediction;instruction fetching;trace cache;instruction cache;superscalar processors

297709

Automatic Compiler-Inserted Prefetching for Pointer-Based Applications.

AbstractAs the disparity between processor and memory speeds continues to grow, memory latency is becoming an increasingly important performance bottleneck. While software-controlled prefetching is an attractive technique for tolerating this latency, its success has been limited thus far to array-based numeric codes. In this paper, we expand the scope of automatic compiler-inserted prefetching to also include the recursive data structures commonly found in pointer-based applications.We propose three compiler-based prefetching schemes, and automate the most widely applicable scheme (greedy prefetching) in an optimizing research compiler. Our experimental results demonstrate that compiler-inserted prefetching can offer significant performance gains on both uniprocessors and large-scale shared-memory multiprocessors.

Introduction OFTWARE -controlled data prefetching [1], [2] offers the potential for bridging the ever-increasing speed gap between the memory subsystem and today's high-performance processors. In recognition of this potential, a number of recent processors have added support for prefetch instructions [3], [4], [5]. While prefetching has enjoyed considerable success in array-based numeric codes [6], its potential in pointer-based applications has remained largely unexplored. This paper investigates compiler-inserted prefetching for pointer-based applications-in par- ticular, those containing recursive data structures. Recursive Data Structures (RDSs) include familiar objects such as linked lists, trees, graphs, etc., where individual nodes are dynamically allocated from the heap, and nodes are linked together through pointers to form the over-all structure. For our purposes, "recursive data structures" can be broadly interpreted to include most pointer-linked data structures (e.g., mutually-recursive data structures, or even a graph of heterogeneous objects). From a memory performance perspective, these pointer-based data structures are expected to be an important concern for the following reasons. For an application to suffer a large memory penalty due to data replacement misses, it typically must have a large data set relative to the cache size. Aside from multi-dimensional arrays, recursive data structures are one of the most common and convenient methods of building large data structures (e.g, B-trees in database applications, octrees in graphics applications, etc. As we traverse a C.-K. Luk is with the Department of Computer Science, University of Toronto, Toronto, Ontario M5S 3G4, Canada. E-mail: [email protected]. T. C. Mowry is with the Computer Science Department, Carnegie Mellon University, Pittsburgh, PA 15213. E-mail: [email protected]. large RDS, we may potentially visit enough intervening nodes to displace a given node from the cache before it is revisited; hence temporal locality may be poor. Finally, in contrast with arrays-where consecutive elements are at contiguous addresses-there is little inherent spatial locality between consecutively-accessed nodes in an RDS, since they are dynamically allocated at arbitrary addresses. To cope with the latency of accessing these pointer-based data structures, we propose three compiler-based schemes for prefetching RDSs, as described in Section II. We implemented the most widely-applicable of these schemes-greedy prefetching-in a modern research compiler (SUIF [7]), as discussed in Section III. To evaluate our schemes, we performed detailed simulations of their impact on both uniprocessor and multiprocessor systems in Sections IV and V, respectively. Finally, we present related work and conclusions in Sections VI and VII. II. Software-Controlled Prefetching for RDSs A key challenge in successfully prefetching RDSs is scheduling the prefetches sufficiently far in advance to fully hide the latency, while introducing minimal runtime overhead. In contrast with array-based codes, where the prefetching distance can be easily controlled using software pipelining [2], the fundamental difficulty with RDSs is that we must first dereference pointers to compute the prefetch addresses. Getting several nodes ahead in an RDS traversal typically involves following a pointer chain. However, the very act of touching these intermediate nodes along the pointer chain means that we cannot tolerate the latency of fetching more than one node ahead. To overcome this pointer-chasing problem [8], we propose three schemes for generating prefetch addresses without following the entire pointer chain. The first two schemes- greedy prefetching and history-pointer prefetching-use a pointer within the current node as the prefetching address; the difference is that greedy prefetching uses existing point- ers, whereas history-pointer prefetching creates new point- ers. The third scheme-data-linearization prefetching- generates prefetch addresses without pointer dereferences. A. Greedy Prefetching In a k-ary RDS, each node contains k pointers to other nodes. Greedy prefetching exploits the fact that when only one of these k neighbors can be immediately followed as the next node in the traversal, but there is often a good chance that other neighbors will be visited sometime in the future. Therefore by prefetching all k pointers when a node is first visited, we hope that enough of these preorder(treeNode * t) f prefetch(t!left); prefetch(t!right); preorder(t!left); preorder(t!right); 4 5partial latency cache miss cache hit cache miss 9 15122 (a) Code with Greedy Prefetching (b) Cache Miss Behavior Fig. 1. Illustration of greedy prefetching. prefetches are successful that we can hide at least some fraction of the miss latency. To illustrate how greedy prefetching works, consider the pre-order traversal of a binary tree (i.e. Figure 1(a) shows the code with greedy prefetching added. Assuming that the computation in process() takes half as long as the cache miss latency L, we would want to prefetch two nodes ahead to fully hide the latency. Figure 1(b) shows the caching behavior of each node. We obviously suffer a full cache miss at the root node (node 1), since there was no opportunity to fetch it ahead of time. However, we would only suffer half of the miss penalty ( L when we visit node 2, and no miss penalty when we eventually visit node 3 (since the time to visit the subtree rooted at node 2 is greater than L). In this example, the latency is fully hidden for roughly half of the nodes, and reduced by 50% for the other half (minus the root node). Greedy prefetching offers the following advantages: (i) it has low runtime overhead, since no additional storage or computation is needed to construct the prefetch point- ers; (ii) it is applicable to a wide variety of RDSs, regardless of how they are accessed or whether their structure is modified frequently; and (iii) it is relatively straightforward to implement in a compiler-in fact, we have implemented it in the SUIF compiler, as we describe later in Section III. The main disadvantage of greedy prefetching is that it does not offer precise control over the prefetching distance, which is the motivation for our next algorithm. B. History-Pointer Prefetching Rather than relying on existing pointers to approximate prefetch addresses, we can potentially synthesize more accurate pointers based on the observed RDS traversal pat- terns. To prefetch d nodes ahead under the history-pointer prefetching scheme [8], we add a new pointer (called a history-pointer) to a node n i to record the observed address of n i+d (the node visited d nodes after n i ) on a recent traversal of the RDS. On subsequent traversals of the RDS, we prefetch the nodes pointed to by these history- pointers. This scheme is most effective when the traversal pattern does not change rapidly over time. To construct the history-pointers, we maintain a FIFO queue of length d which contains pointers to the last d nodes that have just been visited. When we visit a new node n i , the oldest node in the queue will be n i\Gammad (i.e. the node visited d nodes ear- lier), and hence we update the history-pointer of n i\Gammad to point to n i . After the first complete traversal of the RDS, all of the history-pointers will be set. In contrast with greedy prefetching, history-pointer prefetching offers no improvement on the first traversal of an RDS, but can potentially hide all of the latency on subsequent traversals. While history-pointer prefetching offers the potential advantage of improved latency tolerance, this comes at the expense of (i) execution overhead to construct the history-pointers, and (ii) space overhead for storing these new pointers. To minimize execution overhead, we can potentially update the history-pointers less frequently, depending on how rapidly the RDS structure changes. In one extreme, if the RDS never changes, we can set the history-pointers just once. The problem with space overhead is that it potentially worsens the caching behavior. The desire to eliminate this space overhead altogether is the motivation for our next prefetching scheme. C. Data-Linearization Prefetching The idea behind data-linearization prefetching [8] is to map heap-allocated nodes that are likely to be accessed close together in time into contiguous memory locations. With this mapping, one can easily generate prefetch addresses and launch them early enough. Another advantage of this scheme is that it improves spatial locality. The major challenge, however, is how and when we can generate this data layout. In theory, one could dynamically remap the data even after the RDS has been initially constructed, but doing so may result in large runtime overheads and may also violate program semantics. Instead, the easiest time to map the nodes is at creation time, which is appropriate if either the creation order already matches the traversal or- der, or if it can be safely reordered to do so. Since dynamic remapping is expensive (or impossible), this scheme obviously works best if the structure of the RDS changes only slowly (or not at all). If the RDS does change radically, the program will still behave correctly, but prefetching will not improve performance. III. Implementation of Greedy Prefetching Of the three schemes that we propose, greedy prefetching is perhaps the most widely applicable since it does not rely on traversal history information, and it requires no additional storage or computation to construct prefetch ad- dresses. For these reasons, we have implemented a version of greedy prefetching within the SUIF compiler [7], and we will simulate the other two algorithms by hand. Our implementation consists of an analysis phase to recognize RDS accesses, and a scheduling phase to insert prefetches. A. Analysis: Recognizing RDS Accesses To recognize RDS accesses, the compiler uses both type declaration information to recognize which data objects are RDSs, and control structure information to recognize when these objects are being traversed. An RDS type is a record type r containing at least one pointer that points either directly or indirectly to a record type s. (Note that r and s are not restricted to be the same type, since RDSs may struct T f int data; struct T *left; struct T *right; struct A f int struct B *kids[8]; struct C f int j; double f; (a) RDS type (b) RDS type (c) Not RDS type Fig. 2. Examples of which types are recognized as RDS types. while (l) f list *m; for (.) f list *n; f(tree *t) f k(tree tn) f (a) (b) (c) (d) Fig. 3. Examples of control structures recognized as RDS traversals. be comprised of heterogeneous nodes.) For example, the type declarations in Figure 2(a) and Figure 2(b) would be recognized as RDS types, whereas Figure 2(c) would not. After discovering data structures with the appropriate types, the compiler then looks for control structures that are used to traverse the RDSs. In particular, the compiler looks for loops or recursive procedure calls such that during each new loop iteration or procedure invocation, a pointer p to an RDS is assigned a value resulting from a dereference of p-we refer to this as a recurrent pointer update. This heuristic corresponds to how RDS codes are typically written. To detect recurrent pointer updates, the compiler propagates pointer values using a simplified (but less pre- cise) version of earlier pointer analysis algorithms [9], [10]. Figure 3 shows some example program fragments that our compiler treats as RDS accesses. In Figure 3(a), l is updated to l!next!next inside the while-loop. In Figure 3(b), n is assigned the result of the function call g(n) inside the for-loop. (Since our implementation does not perform interprocedural analysis, it assumes that g(n) results in a value n!.!next.) In Figure 3(c), two dereferences of the function argument t are passed as the parameters to two recursive calls. Figure 3(d) is similar to Figure 3(c), except that a record (rather than a pointer) is passed as the function argument. Ideally, the next step would be to analyze data locality across RDS nodes to eliminate unnecessary prefetches. Although we have not automated this step in our compiler, we evaluated its potential benefits in an earlier study [8]. B. Scheduling Prefetches Once RDS accesses have been recognized, the compiler inserts greedy prefetches as follows. At the point where an RDS object is being traversed-i.e. where the recurrent pointer update occurs-the compiler inserts prefetches of all pointers within this object that point to RDS-type objects at the earliest points where these addresses are available within the surrounding loop or procedure body. The availability of prefetch addresses is computed by prop- while (l) f while (l) f prefetch(l!next); (a) Loop tree *q; if (test(t!data)) else if (q != NULL) tree *q; prefetch(t!left); prefetch(t!right); if (test(t!data)) else if (q != NULL) (b) Procedure Fig. 4. Examples of greedy prefetch scheduling. I Benchmark characteristics. Node Recursive Data Input Memory Benchmark Structures Used Data Set Allocated octree Bisort Binary tree 250,000 1,535 KB integers EM3D Singly-linked lists 2000 H-nodes, 1,671 KB 100 E-nodes, 75% local Health Four-way tree and level = 5, 925 KB doubly-linked lists MST Array of singly- 512 nodes 10 KB linked lists Perimeter A quadtree 4Kx4K image 6,445 KB Power Multi-way tree and 10,000 418 KB singly-linked lists customers TreeAdd Binary tree 1024K nodes 12,288 KB Binary tree and 100,000 cities 5,120 KB doubly-linked lists Voronoi Binary tree 20,000 points 10,915 KB agating the earliest generation points of pointer values along with the values themselves. Two examples of greedy prefetch scheduling are shown in Figure 4. Further details of our implementation can be found in Luk's thesis [11]. IV. Prefetching RDSs on Uniprocessors In this section, we quantify the impact of our prefetching schemes on uniprocessor performance. Later, in Section V, we will turn our attention to multiprocessor systems. A. Experimental Framework We performed detailed cycle-by-cycle simulations of the entire Olden benchmark suite [12] on a dynamically- scheduled, superscalar processor similar to the MIPS R10000 [5]. The Olden benchmark suite contains ten pointer-based applications written in C, which are briefly summarized in Table I. The rightmost column in Table I shows the amount of memory dynamically allocated to RDS nodes. Our simulation model varies slightly from the actual MIPS R10000 (e.g., we model two memory units, and we II Uniprocessor simulation parameters. Pipeline Parameters Issue Width 4 Functional Units 2 Int, 2 FP, 2 Memory, 1 Branch Reorder Buffer Size Integer Multiply 12 cycles Integer Divide 76 cycles All Other Integer 1 cycle FP Divide 15 cycles FP Square Root 20 cycles All Other FP 2 cycles Branch Prediction Scheme 2-bit Counters Memory Parameters Primary Instr and Data Caches 16KB, 2-way set-associative Unified Secondary Cache 512KB, 2-way set-associative Line Size 32B Primary-to-Secondary Miss 12 cycles Primary-to-Memory Miss 75 cycles Data Cache Miss Handlers 8 Data Cache Banks 2 Data Cache Fill Time 4 cycles (Requires Exclusive Access) Main Memory Bandwidth 1 access per 20 cycles assume that all functional units are fully-pipelined), but we do model the rich details of the processor including the pipeline, register renaming, the reorder buffer, branch pre- diction, instruction fetching, branching penalties, the memory hierarchy (including contention), etc. Table II shows the parameters of our model. We use pixie [13] to instrument the optimized MIPS object files produced by the com- piler, and pipe the resulting trace into our simulator. To avoid misses during the initialization of dynamically- allocated objects, we used a modified version of the IRIX mallopt routine [14] whereby we prefetch allocated objects before they are initialized. Determining these prefetch addresses is straightforward, since objects of the same size are typically allocated from contiguous memory. This optimization alone led to over twofold speedups relative to using malloc for the majority of the applications- particularly those that frequently allocate small objects. B. Performance of Greedy Prefetching Figure 5 shows the results of our uniprocessor experi- ments. The overall performance improvement offered by greedy prefetching is shown in Figure 5(a), where the two bars correspond to the cases without prefetching (N) and with greedy prefetching (G). These bars represent execution time normalized to the case without prefetching, and they are broken down into four categories explaining what happened during all potential graduation slots. (The number of graduation slots is the issue width-4 in this case- multiplied by the number of cycles.) The bottom section (busy) is the number of slots when instructions actually graduate, the top two sections are any non-graduating slots that are immediately caused by the oldest instruction suffering either a load or store miss, and the inst stall section is all other slots where instructions do not graduate. Note that the load stall and store stall sections are only a first-order approximation of the performance loss due to cache misses, since these delays also exacerbate subsequent data dependence stalls. As we see in Figure 5(a), half of the applications enjoy a speedup ranging from 4% to 45%, and the other half are within 2% of their original performance. For the applications with the largest memory stall penalties-i.e. health, perimeter, and treeadd-much of this stall time has been eliminated. In the cases of bisort and mst, prefetching overhead more than offset the reduction in memory stalls (thus resulting in a slight performance degradation), but this was not a problem in the other eight applications. To understand the performance results in greater depth, Figure breaks down the original primary cache misses into three categories: (i) those that are prefetched and subsequently hit in the primary cache (pf hit), (ii) those that are prefetched but remain primary misses (pf miss), and (iii) those that are not prefetched (nopf miss). The sum of the pf hit and pf miss cases is also known as the coverage factor, which ideally should be 100%. For em3d, power, and voronoi, the coverage factor is quite low (un- der 20%) because most of their misses are caused by array or scalar references-hence prefetching RDSs yields little improvement. In all other cases, the coverage factor is above 60%, and in four cases we achieve nearly perfect coverage. If the pf miss category is large, this indicates that prefetches were not scheduled effectively-either they were issued too late to hide the latency, or else they were too early and the prefetched data was displaced from the cache before it could be referenced. This category is most prominent in mst, where the compiler is unable to prefetch early enough during the traversal of very short linked lists within a hash table. Since greedy prefetching offer little control over prefetching distance, it is not surprising that scheduling is imperfect-in fact, it is encouraging that the pf miss fractions are this low. To help evaluate the costs of prefetching, Figure 5(c) shows the fraction of dynamic prefetches that are unnecessary because the data is found in the primary cache. For each application, we show four different bars indicating the total (dynamic) unnecessary prefetches caused by static prefetch instructions with hit rates up to a given threshold. Hence the bar labeled "100" corresponds to all unnecessary prefetches, whereas the bar labeled "99" shows the total unnecessary prefetches if we exclude prefetch instructions with hit rates over 99%, etc. This breakdown indicates the potential for reducing overhead by eliminating static prefetch instructions that are clearly of little value. For example, eliminating prefetches with hit rates over 99% would eliminate over half of the unnecessary prefetches in perimeter, thus decreasing overhead significantly. In con- trast, reducing overhead with a flat distribution (e.g., bh) is more difficult since prefetches that sometimes hit also miss at least 10% of the time; therefore, eliminating them may sacrifice some latency-hiding benefit. We found that eliminating prefetches with hit rates above 95% improves performance by 1-7% for these applications [8]. Finally, we measured the impact of greedy prefetching on memory bandwidth consumption. We observe that on av- 0Normalized Execution Time load stall 100.0 96.6 100.0 101.2 100.0 99.8 100.0100.0 101.6 100.0100.0 99.9 100.0100.0100.0 99.4 bh bisort em3d health mst perimeter power treeadd tsp voronoi store stall inst stall busy (a) Execution Time |||||||% of Original Load D-Cache nopf_miss bisort health perimeter treeadd voronoi bh em3d mst power tsp pf_miss pf_hit |||% of that Hit in 10099 95 90 10099 95 90 10099 95 90 10099 95 90 10099 95 90 10099 95 90 10099 95 90 10099 95 90 10099 95 90 10099 95 90 bh bisort em3d health mst perimeter power treeadd tsp voronoi (b) Coverage Factor (c) Unnecessary Prefetches Fig. 5. Performance impact of compiler-inserted greedy prefetching on a uniprocessor. erage, greedy prefetching increases the traffic between the primary and secondary caches by 12.7%, and the traffic between the secondary cache and main memory by 7.8%. In our experiments, this has almost no impact on perfor- mance. Hence greedy prefetching does not appear to be suffering from memory bandwidth problems. In summary, we have seen that automatic compiler-inserted prefetching can result in significant speedups for uniprocessor applications containing RDSs. We now investigate whether the two more sophisticated prefetching schemes can offer even larger performance gains. C. Performance of History-Pointer Prefetching and Data- Linearization Prefetching We applied history-pointer prefetching and data- linearization prefetching by hand to several of our applica- tions. History-pointer prefetching is applicable to health because the list structures that are accessed by a key procedure remain unchanged across the over ten thousand times that it is called. As a result, history-pointer prefetching achieves a 40% speedup over greedy prefetching through better miss coverage and fewer unnecessary prefetches. Although history-pointer prefetching has fewer unnecessary prefetches than greedy prefetching, it has significantly higher instruction overhead due to the extra work required to maintain the history-pointers. Data-linearization prefetching is applicable to both perimeter and treeadd, because the creation order is identical to the major subsequent traversal order in both cases. As a result, data linearization does not require changing the data layout in these cases (hence spatial locality is unaffected). By reducing the number of unnecessary prefetches (and hence prefetching overhead) while maintaining good coverage factors, data-linearization prefetching results in speedups of 9% and 18% over greedy prefetching for perimeter and treeadd, respectively. Overall, we see that both schemes can potentially offer significant improvements over greedy prefetching when applicable. V. Prefetching RDSs on Multiprocessors Having observed the benefits of automatic prefetching of RDSs on uniprocessors, we now investigate whether the compiler can also accelerate pointer-based applications running on multiprocessors. In earlier studies, Mowry demonstrated that the compiler can successfully prefetch parallel matrix-based codes [2], [15], but the compiler used in those studies did not attempt to prefetch pointer-based access patterns. However, through hand-inserted prefetch- ing, Mowry was able to achieve a significant speedup in BARNES [15], which is a pointer-intensive shared-memory parallel application from the SPLASH suite [16]. BARNES performs a hierarchical n-body simulation of the evolution of galaxies. The main computation consists of a depth-first traversal of an octree structure to compute the gravitational force exerted by the given body on all other bodies in the tree. This is repeated for each body in the system, and the bodies are statically assigned to processors for the duration of each time step. Cache misses occur whenever a processor visits a part of the octree that is not already in its cache, either due to replacements or communication. To insert prefetches by hand, Mowry used a strategy similar to greedy prefetching: upon first arriving at a node, he prefetched all immediate children before descending depth-first into the first child. III Memory latencies in multiprocessor simulations. Destination of Access Read Write Primary Cache 1 cycle 1 cycle Secondary Cache 15 cycles 4 cycles Remote Node 101 cycles 89 cycles Dirty Remote, Remote Home 132 cycles 120 cycles ||||||||Normalized Execution Time memory stalls synchronization instructions86 85 of Original D-Cache nopf_miss pf_miss pf_hit ||||||% of that Hit in D-Cache (a) Execution (b) Coverage (c) Unnecessary Time Factor Prefetches Fig. 6. Impact of compiler-inserted greeding prefetching on BARNES on a multiprocessor compiler-inserted greedy prefetching, hand-inserted prefetching). To evaluate the performance of our compiler-based implementation of greedy prefetching on a multiprocessor, we compared it with hand-inserted prefetching for BARNES. For the sake of comparison, we adopted the same simulation environment used in Mowry's earlier study [15], which we now briefly summarize. We simulated a cache-coherent, shared-memory multiprocessor that resembles the DASH multiprocessor [17]. Our simulated machine consists of 16 processors, each of which has two levels of direct-mapped caches, both using 16 byte lines. Table III shows the latency for servicing an access to different levels of the memory hierarchy, in the absence of contention (our simulations did model contention, however). To make simulations fea- sible, we scaled down both the problem size and cache sizes accordingly (we ran 8192 bodies through 3 times steps on an 8K/64K cache hierarchy), as was done (and explained in more detail) in the original study [2]. Figure 6 shows the impact of both compiler-inserted greedy prefetching (G) and hand-inserted prefetching (H) on BARNES. The execution times in Figure 6(a) are broken down as follows: the bottom section is the amount of time spent executing instructions (including any prefetching instruction overhead), and the middle and top sections are synchronization and memory stall times, respectively. As we see in Figure 6(a), the compiler achieves nearly identical performance to hand-inserted prefetching. The compiler prefetches 90% of the original cache misses with only 15% of these misses being unnecessary, as we see in Figures 6(b) and 6(c), respectively. Of the prefetched misses, the latency was fully hidden in half of the cases (pf hit), and partially hidden in the other cases (pf miss). By eliminating roughly half of the original memory stall time, the compiler was able to achieve a 16% speedup. The compiler's greedy strategy for inserting prefetches is quite similar to what was done by hand, with the following exception. In an effort to minimize unnecessary prefetches, the compiler's default strategy is to prefetch only the first 64 bytes within a given RDS node. In the case of BARNES, the nodes are longer than 64 bytes, and we discovered that hand-inserted prefetching achieves better performance when we prefetch the entire nodes. In this case, the improved miss coverage of prefetching the entire nodes is worth the additional unnecessary prefetches, thereby resulting in a 1% speedup over compiler-inserted prefetching. Overall, however, we are quite pleased that the compiler was able to do this well, nearly matching the best performance that we could achieve by hand. VI. Related Work Although prefetching has been studied extensively for array-based numeric codes [6], [18], relatively little work has been done on non-numeric applications. Chen et al. [19] used global instruction scheduling techniques to move address generation back as early as possible to hide a small cache miss latency (10 cycles), and found mixed results. In contrast, our algorithms focus only on RDS accesses, and can issue prefetches much earlier (across procedure and loop iteration boundaries) by overcoming the pointer-chasing problem. Zhang and Torrellas [20] proposed a hardware-assisted scheme for prefetching irregular applications in shared-memory multiprocessors. Under their scheme, programs are annotated to bind together groups of data (e.g., fields in a record or two records linked by a pointer), which are then prefetched under hardware con- trol. Compared with our compiler-based approach, their scheme has two shortcomings: (i) annotations are inserted manually, and (ii) their hardware extensions are not likely to be applicable in uniprocessors. Joseph and Grunwald [21] proposed a hardware-based Markov prefetching scheme which prefetches multiple predicted addresses upon a primary cache miss. While Markov prefetching can potentially handle chaotic miss patterns, it requires considerably more hardware support and has less flexibility in selecting what to prefetch and controlling the prefetch distance than our compiler-based schemes. To our knowledge, the only compiler-based pointer prefetching scheme in the literature is the SPAID scheme proposed by Lipasti et al. [22]. Based on an observation that procedures are likely to dereference any pointers passed to them as arguments, SPAID inserts prefetches for the objects pointed to by these pointer arguments at the call sites. Therefore this scheme is only effective if the interval between the start of a procedure call and its dereference of a pointer is comparable to the cache miss latency. In an earlier study [8], we found that greedy prefetching offers substantially better performance than SPAID by hiding more latency while paying less overhead. VII. Conclusions While automatic compiler-inserted prefetching has shown considerable success in hiding the memory latency of array-based codes, the compiler technology for successfully prefetching pointer-based data structures has thus far been lacking. In this paper, we propose three prefetching schemes which overcome the pointer-chasing problem, we automate the most widely applicable scheme (greedy prefetching) in the compiler, and we evaluate its performance on both a modern superscalar uniprocessor (sim- ilar to the MIPS R10000) and on a large-scale shared-memory multiprocessor. Our uniprocessor experiments show that automatic compiler-inserted prefetching can accelerate pointer-based applications by as much as 45%. In addition, the more sophisticated algorithms (which we currently simulate by hand) can offer even larger performance gains. Our multiprocessor experiments demonstrate that the compiler can potentially provide equivalent performance to hand-inserted prefetching even on parallel ap- plications. These encouraging results suggest that the latency problem for pointer-based codes may be addressed largely through the prefetch instructions that already exist in many recent microprocessors. Acknowledgments This work is supported by a grant from IBM Canada's Centre for Advanced Studies. Chi-Keung Luk is partially supported by a Canadian CommonwealthFellowship. Todd C. Mowry is partially supported by a Faculty Development Award from IBM. --R "Software prefetching," Tolerating Latency Through Software-Controlled Data Prefetching "Com- piler techniques for data prefetching on the PowerPC," "Data prefetching on the HP PA8000," "The MIPS R10000 superscalar microprocessor," "Design and evaluation of a compiler algorithm for prefetching," "SUIF: An infrastructure for research on parallelizing and optimizing compilers," "Compiler-based prefetching for recursive data structures," "Context-sensitive interprocedural points-to analysis in the presence of function pointers," "Interprocedural modification side effect analysis with pointer aliasing," Optimizing the Cache Performance of Non-Numeric Applications "Support- ing dynamic data structures on distributed memory machines," "Tracing with pixie," "Fast fits," "Tolerating latency in multiprocessors through compiler-inserted prefetching," parallel applications for shared memory," "The Stanford DASH multiproces- sor," "An effective on-chip preloading scheme to reduce data access penalty," "Data access microarchitectures for superscalar processors with compiler-assisted data prefetching," "Speeding up irregular applications in shared-memory multiprocessors: Memory binding and group prefetching," "Prefetching using Markov predic- tors," "SPAID: Software prefetching in pointer- and call-intensive environments," --TR --CTR Subramanian Ramaswamy , Jaswanth Sreeram , Sudhakar Yalamanchili , Krishna V. Palem, Data trace cache: an application specific cache architecture, ACM SIGARCH Computer Architecture News, v.34 n.1, March 2006 Shimin Chen , Phillip B. Gibbons , Todd C. Mowry, Improving index performance through prefetching, ACM SIGMOD Record, v.30 n.2, p.235-246, June 2001 Tatsushi Inagaki , Tamiya Onodera , Hideaki Komatsu , Toshio Nakatani, Stride prefetching by dynamically inspecting objects, ACM SIGPLAN Notices, v.38 n.5, May Evangelia Athanasaki , Nikos Anastopoulos , Kornilios Kourtis , Nectarios Koziris, Exploring the performance limits of simultaneous multithreading for memory intensive applications, The Journal of Supercomputing, v.44 n.1, p.64-97, April 2008 Chi-Keung Luk, Tolerating memory latency through software-controlled pre-execution in simultaneous multithreading processors, ACM SIGARCH Computer Architecture News, v.29 n.2, p.40-51, May 2001

compiler optimization;prefetching;performance evaluation;caches;shared-memory multiprocessors;recursive data structures;pointer-based applications

297713

Analysis of Temporal-Based Program Behavior for Improved Instruction Cache Performance.

AbstractIn this paper, we examine temporal-based program interaction in order to improve layout by reducing the probability that program units will conflict in an instruction cache. In that context, we present two profile-guided procedure reordering algorithms. Both techniques use cache line coloring to arrive at a final program layout and target the elimination of first generation cache conflicts (i.e., conflicts between caller/callee pairs). The first algorithm builds a call graph that records local temporal interaction between procedures. We will describe how the call graph is used to guide the placement step and present methods that accelerate cache line coloring by exploring aggressive graph pruning techniques. In the second approach, we capture global temporal program interaction by constructing a Conflict Miss Graph (CMG). The CMG estimates the worst-case number of misses two competing procedures can inflict upon one another and reducing higher generation cache conflicts. We use a pruned CMG graph to guide cache line coloring. Using several C and C++ benchmarks, we show the benefits of letting both types of graphs guide procedure reordering to improve instruction cache hit rates. To contrast the differences between these two forms of temporal interaction, we also develop new characterization streams based on the Inter-Reference Gap (IRG) model.

Introduction C ACHE memories are found on most microprocessors designed today. Caching the instruction stream can be very beneficial since instruction references exhibit a high degree of spatial and temporal locality. Still, cache misses will occur for one of three reasons [1]: 1. first time reference, 2. finite cache capacity, or 3. memory address conflict. Our work here is focused on reducing memory address conflicts by rearranging a program on the available memory space. Analysis of program interaction can be performed at a range of granularities, the coarsest being an individual procedure [2]. We begin by considering the procedure Call Graph Ordering (CGO) associated with a program. The CGO captures local temporal interaction by weighting its edges with the number of times one procedure follows another during program execution. We also consider the interaction of basic blocks contained within procedures by identifying the number of cache lines touched by each basic block in our Conflict Miss Graph (CMG). We do not attempt to move basic blocks or split procedures [3] though. We weight CMG edges by measuring global temporal interaction between procedures occurring in a finite window, containing as many entries as there are cache lines. Program interaction outside this window is not of interest because of the finite cache effect. We use these graphs as input to our coloring algorithm to produce an improved code layout for instruction caches. To characterize the temporal behavior captured by these graphs, we extend the Inter-Reference Gap (IRG) model [4]. We define three new IRG-based streams that describe different levels of procedure-based temporal interaction. We show how we can use them to compare the temporal content between the CGO, CMG and the Temporal Relationship Graph (TRG) (as described by Gloy et al. in [5]). There has been a considerable amount of work done on code repositioning for improved instruction cache performance [3], [5], [6], [7], [8], [9], [10]. In the following section we discuss some of this work, as it relates to our work here. A. Related Work Pettis and Hansen [3] employ procedure and basic block reordering as well as procedure splitting based on frequency counts to minimize instruction cache conflicts. The layout of a program is directed by traversing call graph edges in decreasing edge weight order using a closest-is-best placement strategy. Chains are formed by merging nodes, laying them out next to each other until the entire graph is processed A number of related techniques have been proposed, focusing on mapping loops [6], operating system code [10], traces [9], and activity sets [7]. Two other approaches discussed in [8] and [11] reorganize code based on compile-time information. The profile-guided algorithms described above use calling frequencies to weight a graph and guide placement [3], [6], [9], [10]. Our first approach also uses calling frequencies but improves performance by intelligently placing procedures in the cache by coloring cache lines. The second algorithm described in this paper captures global temporal information and attempts to minimize conflicts present between procedures that do not immediately follow each other during execution. It is similar in spirit with the approach described in [5], with some differences that will be high-lighted in Section V. Also, our graph coloring algorithm works at a finer level of granularity (cache line size instead of cache size [6], [11]), and can avoid conflicts encountered when either forming chains with the closest-is-best heuristic [3] or dealing with subgraphs having a size larger than the cache. This paper is organized as follows. In Sections II and III we describe our graph construction algorithms. Section II describes an improved graph pruning technique. In Section IV we report simulation results. Section V reviews the IRG temporal analysis model and presents new methods for characterizing program interaction. II. Call Graph Ordering Program layout may involve two steps: 1) constructing a graph-based representation of the program and 2) using the graph to perform layout of the program on the available memory space. A Call Graph is a procedure graph having edges between procedures that call each other. The edges are weighted with the call/return frequency captured in the program profile. Each procedure is mapped to a single vertex, with all call paths between any two procedures condensed into a single edge between the two vertices in the graph. Edge weights can be derived from profiling information or estimated from the program control flow [8], [12]. In this paper we concentrate on profile-based edge weights. After constructing the Call Graph we lay out the program using cache line coloring. We start by dividing the cache into a set of colors, one color for each cache line. For each procedure, we count the number of cache lines needed to hold the procedure, record the cache colors used to map the procedure, and keep track of the unavailable-set of colors (i.e., the cache lines where the procedure should not be mapped to). We define the popular procedure set as those procedures which are frequently visited. The popular edge set contains the frequently traversed edges. The rest of the procedures (and edges) will be called unpopular. Unpopular procedures are pruned from the graph. Pruning reduces the amount of work for placement, and allows us to focus on the procedures most likely to encounter misses. A discussion of the base pruning algorithm used can be found in [2]. Note, there is a difference between popular procedures and procedures that consume a noticeable portion of a pro- gram's overall execution time. A time consuming procedure may be labeled unpopular because it rarely switches control flow to another procedure. If a procedure rarely switches control flow, it causes a small number of conflicts misses with the rest of procedures. The algorithm sorts the popular edges in descending edge weight order. We then traverse the sorted popular edge list, inspect the state (i.e., mapped or unmapped) of the two procedures forming the edge and map the procedures using heuristics. Figure 1 provides a pseudo code description of color mapping. A more complete description can be found in [2]. This process is repeated until all of the edges in the popular set have been processed. The unpopular procedures fill the holes left from coloring using a simple Input: graph G(P,E), descending order based on weight; Eliminate all threshold T; FOR-EACH (remaining edge between procedures Pi and Pj) - SELECT(state of Pi, state of Pj) CASE I: (Pi is unmapped and Pj is unmapped) - Arbitrarily map Pi and Pj, forming a compound node Pij CASE II: (both Pi and Pj are mapped, but they reside in two different compound nodes Ci and Cj) - Concatenate the two compound nodes Ci and Cj, minimizing the distance between Pi and Pj; If there is a color conflict, shift Ci in the color space until there is no conflict; If a conflict can not be avoided, return Ci to its original position; CASE III: (either Pi or Pj is mapped, but not both) - same as CASE II CASE IV: (both Pi and Pj are mapped and they belong to the same compound node Ck) - If there is no conflict, return them in position; Else - Move the procedure closest to an end of the compound node Ck (Pi) to that end of Ck (outside of the compound node); If there is still a conflict, shift Pi in the color space until no conflict occurs with Pj; If conflict can not be avoided, leave Pi to its original position; Update the unavailable set of colors; Fill in holes created by CASES II, III, and IV; Fig. 1. Pseudo code for our cache line coloring algorithm. depth-first traversal of the unpopular edges joining them. The algorithm in Fig.1 assumes a direct-mapped cache organization. For associative caches our algorithm breaks up the address space into chunks, equal in size to (number of cache sets cache line size). Therefore, the number of sets represents the number of available colors in the mapping. The modified color mapping algorithm keeps track of the number of times each color (set) appears in the procedure's unavailable-set of colors. Mapping a procedure to a color (set) does not cause any conflicts as long as the number of times that color (set) appears in the unavailable- set of colors is less than the degree of associativity of the cache. Next, we look at how to efficiently eliminate a majority of the work spent on coloring by using an aggressive graph pruning algorithm. A. Pruning Rules For Procedure Call Graphs Pruning a call graph is done using a fixed threshold value (selected edge weight) [2] . In this section we present pruning rules that can reduce the size of the graph that is used in cache coloring. They are specifically designed to reduce number of lines in the cache; number of procedures P (nodes) in graph G; Sum of all incident edges on procedure Pi in Graph G; Sort procedures based on increasing NodeWeight DO-WHILE (at least one procedure is pruned) - FOR-EACH unpruned procedure Pi - number of neighbors of Pi; number of cache lines comprising Pi; sum of the sizes of the num-i neighbors of Pi; Prune procedure Pi and all edges incident on Pi; Fig. 2. Pseudo code for our C pruning rule. the number of first-generation cache conflicts. We assume we are using a direct-mapped cache containing C cache lines. The program is represented as an undirected graph (P,E) where nodes and each edge (i; represents a procedure call in the program. The number of cache lines spanned by each procedure is size i . For each edge (i; j), weight[i; j] is the number of times procedures i and j follow one another in the control flow (in either order). A procedure mapping M is an assignment of each procedure i to size[i] adjacent cache lines within the cache (with wraparound). The cost of a procedure mapping is the sum of all weight[i; j] for all procedures i and j, such that (i; overlap in the cache. An optimal mapping is one that is less costly than any other mapping. Note that the cost of a mapping depends only on the number of immediately adjacent procedures whose mappings in the cache conflict. Conflicts between procedures that do not call one another are not considered. Further- more, the cost of assigning two adjacent procedures i and j to conflicting cache lines is a constant, equal to the number of conflicts, even though the actual number of replaced cache lines may be smaller. B. C Pruning Rule Consider a cache mapping problem P. It is possible to determine in some cases that a particular node i will be able to be mapped to the cache without causing any con- flicts, regardless of where in the cache all adjacent nodes are eventually mapped. In this case i, and all edges connected to i, can be deleted from the graph, creating a new cache mapping problem P 0 with one less node. Figure 2 provides pseudo code for our pruning algorithm. This pruning rule is a generalization of the rules described in [13] to perform graph coloring. The graph coloring problem is to assign one of K colors to the nodes of the graph such that adjacent nodes are not assigned the same color. If a node has neighbors it can be deleted because, regardless of how its neighbors are eventually col- ored, there will definitely be at least one color left over that can be assigned to it. The deleted nodes are then colored in the reverse order of their deletion. The remaining (non-prunable) graph is passed to our coloring algorithm. Once coloring has been performed, each pruned node must be mapped. The nodes are laid out in the opposite order of their deletion. III. Conflict Miss Graphs Next we consider cache misses which can occur between procedures many procedures away in the call graph, as well as on different call chains [14]. We capture temporal information by weighting the edges of a procedure graph with an estimation of the worst case number of conflict misses that can occur between any two procedures. We then use the graph to apply cache line coloring to place procedures in the cache address space. We call this graph a Conflict Miss Graph (CMG). The complete algorithm is described in [14]. We summarize it here and will contrast it with the CGO in Section V using Inter-Reference Gap analysis. A. Conflict Miss Graph Construction The CMG is built using profile data. We assume a worst-case scenario where procedures completely overlap in the cache address space every time they interact. Given a cache configuration we determine the size of a procedure P i in cache lines. We also compute the number of unique cache lines spanned by every basic block executed by a procedure, l i . We identify the first time a basic block is executed, and label those references as globally unique accesses, gl i The CMG is an undirected procedure graph with edges being weighted according to our worst case miss model [14]. The edge weights are updated based on the contents of an N-entry table, where N is the number of cache lines. The table is fully-associative and uses an LRU replacement pol- icy. Every entry (i.e., cache line) in the table is called live. A procedure that has at least one live cache line is also called live. When P i is activated, we update the edge weights between P i and all procedures that have at least one live cache line and were activated since the last activation of (if this last activation is captured in the LRU table). The LRU table allows us to estimate the finite cache effect. We increment the CMG edge weight between P i and a live procedure P j by the minimum of: (i) the accumulated number of unique live cache lines of P last oc- currence) and (ii) the number of unique cache lines of P i 's current activation (excluding cold-start misses). A detailed example of updating CMG edge weights can be found in [14]. CMG edge weights are more accurate than CGO edge weights because (i) CGO edge weights do not record the number of cache lines that may conflict per call, and (ii) interaction between procedures that do not directly call each other is not captured. IV. Experimental Results We use trace-driven simulation to quantify the instruction cache performance of the resulting layouts. Traces are generated using ATOM [15] on a DEC 3000 AXP workstation running Digital Unix V4.0. All applications are compiled with the DEC C V5.2 and DEC C++ V5.5 com- pilers. The same input is used to train the algorithm and gather performance results. We simulated an 8KB, direct-mapped, instruction cache with a 32-byte line size (similar in design to the DEC Alpha 21064 and 21164 instruction caches). Our benchmark suite includes perl from SPECINT95, flex (generator of lexical analyzers), gs (ghostscript postscript viewer) and bison is a C parser gen- erator. It also includes PC++2dep (C++ front-end written in C/C++), f2dep (Fortran front-end written in C/C++), dep2C++ (C/C++ program translating Sage internal representation to C++ code) and ixx (IDL parser written in C++). Table I presents the static and dynamic characteristics of the benchmarks. Column 2 shows the input used to both test and train our algorithms. Columns 3-5 list the total number of instructions executed, the static size of the application in kilobytes and the number of static procedures in the program. Column 6 presents the percentage of the program that contains popular procedures in the CMG while column 7 contains the percentage of procedures that were found to be popular (for CMG). The last column presents the percentage of unactivated procedures used to fill in the gaps left from the color mapping. To prune the CMG graph, we form the popular set from those procedures that are connected by edges that contribute up to 80% of the total sum of edge weights in the CMG [14]. Notice that the pruning algorithm reduces the size of the CMG by 80-97%, and reduces the size of the executable by 77.7-94.5% of the executable. This allows us to concentrate on the important procedures in the program. A. Simulation Results We compare simulation results against the ordering produced by the DEC compiler (static DFS ordering of proce- dures) and CGO using a fixed threshold value for pruning (no aggressive pruning was employed). Table II shows the instruction cache miss rates. In all cases, the same inputs were used for both training and testing. The first column denotes the application while columns 2-4 (7-9) shows the instruction cache miss rates (number of cache misses) for DFS, CGO and CMG respec- tively. Columns 5 and 6 show the relative improvement of CMG over DFS and CGO respectively. As we can see from Table II, the average instruction cache miss rate for CMG is reduced by 30% on average over the DFS ordering, and by 21% on average over the CGO ordering. CMG improves performance against both static DFS and CGO over all benchmarks except bison, flex and gs. Bison and flex already have a very low miss rate and no further improvement can be achieved. Gs has a large number of popular procedures that can not be mapped in the cache with significant reduction of the miss rate. Next, we apply the pruning C pruning rule to CGO for four benchmarks, bison, flex, gs, and perl. We have also tried to apply this approach to CMG, but found we were unable to significantly reduce the size of the graph. As shown in Tables III and IV, the pruning rule deletes all 125 nodes from bison and completely eliminates all first-generation conflict misses. Similarly, most of the nodes are pruned from the other benchmarks, accompanied by a significant drop in the number of first-generation conflict misses. However, this drop is most of the times followed by an increase in the total number of misses. By deleting nodes and edges that do not contribute to first-generation conflict misses, the coloring algorithm is deprived of information that can be used to prevent higher order misses. In the case of the bison benchmark, the nodes are inserted into the mapping with complete disregard for higher order conflicts. These results suggest that node pruning rules such as C can be useful as part of a cache conflict reduction strategy, but only when paired with other techniques that prevent higher order cache conflicts from canceling out the benefits of reducing first-order conflicts. B. Input training sensitivity Since our procedure reordering algorithm is profile- driven, we tried different training and test input files as shown in Table V. Column 2 (3) has the training (test) input while column 4 shows the size of the traces in millions of instructions for the test and the training inputs (the last one in parentheses). The last three columns present the miss ratios for each of the algorithms simulated. As we can see from Table V, although the performance of both the CGO and the CMG approach drops compared to the simulations using the same inputs, the relative advantage of CMG against CGO and static DFS still remains. In fact the performance gain is of the same order for all benchmarks, i.e. CGO and CMG achieve similar performance for bison and gs, while CMG improves significantly the miss ratio of ixx. V. Temporal locality and procedure-based IRG Next we characterize the temporal interaction exposed by CGO, CMG and the Temporal Relationship Graph using an extended version of the Inter-Reference Gap (IRG) model [4]. A. Procedure-based IRG In [4], Phalke and Gopinath define the IRG for an address as the number of memory references between successive references to that address. An IRG stream for an address in a trace is the sequence of successive IRG values for that address and can be used to characterize its temporal locality. Similarly, we can measure the temporal locality of larger program granules such as basic blocks, cache lines or procedures. The accuracy of the newly generated IRG stream depends on the interval granularity. In this work we set the program unit under study to be a procedure while we vary the interval definition. Program Input Instr. Exe Size # Static Pop Procs Unpop Procs in million in KB Procs % Exe Size % Procs % Exe Size perl primes 12 512 671 4.9 (25.1K) 5.2 3.5 (18.1K) flex fixit 19 112 170 14.8 (16.6K) 17.6 6.5 (7.3K) bison objparse 56 112 158 22.4 (25.1K) 22.1 6.1 (6.9K) pC++2dep sample 19 480 665 9.5 (45.7K) 16.3 10.3 (49.5K) dep2C++ sample 31 560 1338 4.8 (27.1K) 1.7 1.7 (9.5K) gs tiger 34 496 1410 12.9 (64.0K) 11.2 9.3 (46.3K) ixx layout 48 472 1581 5.7 (27.2K) 5.1 2.4 (11.7K) I Attributes of traced applications. The attributes include the number of executed instructions, the application executable size, the number of static procedures, the percentage of program's size occupied by popular procedures, the percentage of procedures that were found to be popular and the percentage of unactivated procedures that were used to fill memory gaps left after applying coloring. I-Cache Miss Rate Reduction # I-Cache Misses Program DFS CGO CMG DFS CGO DFS CGO CMG perl 4.72% 4.60% 3.77% .95 .83 588,123 572,650 469,329 flex 0.53% 0.45% 0.45% .08 .00 100,488 85,538 85,478 bison 0.04% 0.04% 0.05% -.01 -.01 21,798 21,379 26,792 pC++2dep 4.72% 5.46% 3.68% 1.04 1.78 895,261 1,035,639 698,003 dep2C++ 3.92% 3.46% 3.11% .81 .35 1,205,076 1,063,682 957,102 gs 3.45% 2.09% 2.08% 1.37 .01 1,176,335 712,230 709,643 ixx 5.83% 4.42% 2.57% 3.26 1.85 2,843,330 2,154,747 1,251,022 Avg. 3.52% 3.15% 2.48% II Instruction cache performance for static DFS, CGO and CMG-based reordering. Column 1 lists the application. Columns 2-4 show the instruction miss rates. The next two show the percent improvement over each by our algorithm. The last three columns show the number of instruction cache misses. Program Input Visited Pruned Pruned Pruned Procs. 1st pass 2nd pass 3rd pass bison objparse 125 125 0 0 flex fixit 97 61 5 2 gs tiger 532 524 5 0 perl primes 209 113 4 0 III The results of applying the C pruning rules to Call Graphs for four applications. Passes refer to pruning iterations over the graph. The algorithm is finished when no more nodes can be pruned in the graph. Program CGO first higher C* first higher Miss Rate order order Miss Rate order order bison .04 1316 19812 .14 0 79888 flex .45 55004 30280 .51 42467 55233 gs 2.09 530908 658066 2.39 25663 791567 perl 4.60 99327 473070 4.45 92914 462056 IV The results of applying the C pruning rules to Call Graphs for four applications. Cache parameters are the same as those used above. Program Training Test Trace Static DFS CGO CMG input input instr. Miss rate Miss rate Miss rate bison objparse cparse 35.6M (56.1) 0.05% 0.04% 0.06% flex fixit scan 23.2M (19.1) 0.49% 0.43% 0.38% gs tiger golfer 15.5M (34.1) 3.39% 2.49% 2.51% ixx layout widget 52.7M (48.7) 5.89% 4.45% 2.54% perl primes jumble 71.9M (12.4) 4.36% 4.61% 4.16% Miss ratios when using different test and training inputs. The inputs are described in columns 2-3 while the sizes of their corresponding traces are presented in column 4. Columns 5-7 include the instruction cache miss ratios. The original IRG model exploits the temporal locality of a single procedure, but not the temporal interaction between procedures. Therefore, we redefine the IRG value for a procedure pair A; B as the number of unique activated procedures between invocations of A and B. We will refer to this value as the Inter-Reference Procedure Gap (IRPG). The CGO edge weights record part of the IRPG stream since they capture the IRPG values of length 1. In the TRG every node represents a procedure and every edge is weighted by the number of times procedure A vice versa only when both of them are found inside a moving time window. The window includes previously invoked procedures and its size is proportional to the size of the cache. The window's contents are managed as a LRU queue. The temporal interaction recorded by a TRG can be characterized by the Inter-Reference Intermediate Line Gap (IRILG) whose elements are equal to the number of unique cache lines activated between successive A and B invocations. The decision of when to update a TRG edge depends on the size of the window, or equivalently on the values present in the IRILG stream. The edge weight is simply the count of all IRILG elements with a value less than the predefined window size. The TRG captures temporal interaction at a more detailed level than CGO because the IRILG stream is richer in content than the IRPG stream. The CMG edge weight between procedures A and B is updated only when A and B follow one another inside a moving time window proportional in size to the cache size. A CMG edge weight is updated whenever the IRILG value is less than the window size. In both the TRG and the CMG, procedures interact as long as they are found inside the time window. A CMG, however, replaces procedures in the time window based on the age of individual lines in a procedure [14], while a TRG manages replacement on an entire procedure basis [5]. In addition, a CMG edge weight between A and B is incremented by the minimum of the unique live cache lines of the successive invocations of A and B. The TRG simply counts the number of times A and B follow each other. We define the Inter-Reference Active Line Set (IRALS) for a procedure pair as the sequence of the number of unique live cache lines referenced between any successive occurrences of A and B. Each IRALS element value is computed according to the Worst Case Miss analysis presented in Section III. A CMG edge weight is equal to the sum of the IRALS values whose corresponding IRILG values are less than the window size. Table VI shows the IRPG, IRILG, and IRALS element frequency distribution for two edges in the ixx benchmark. The selected edges have the 4th and 12th highest calling frequency in the CGO popular edge list and are labeled as e 4 and e 12 respectively. We classify the stream values in the ranges shown in columns 1,4 and 7, and present per stream frequency distributions in columns 2-3, 5-6 and 8- 9. The numbers in parentheses indicate how closely CGO approximates the temporal information captured by the stream under consideration. For example, while 66.1% of IRILG elements for e 12 lie in the range between 2 and 10 unique cache lines, only 62.4% of them are recorded in the CGO. The 3 different approaches to edge weighting have significant impact on the global edge ordering and the final procedure placement. Ixx CMG and CGO popular edge set intersection Edge index in CGO ordering Edge index in ordering Fig. 3. Relative edge ordering for the intersection of the CMG and CGO popular edge orderings in ixx. Fig. 3 compares the ordering of the common popular edges of CGO and CMG, in ixx. A point at location (3,4) means that the popular edge under consideration was found 3rd in CGO and 4th in CMG ordering. Points on the straight line correspond to edges with the same relative position in both edge lists. Points lying above (below) the straight line indicate edges with a higher priority in the CGO (CMG) edge list. Notice that very few edges fall below the straight line due to the artificially inflated edge index in the CMG edge list (which is much larger than IRPG 4th 12th IRILG 4th 12th IRALS 4th 12th value VI Frequency distribution of the IRPG, IRILG and IRALS sequences for two procedure pairs (4th and 12th in calling frequency) of the ixx benchmark. Program CGO pop CMG pop CGO " CMG CGO pop CMG pop CGO " CMG Procs Procs Procs Edges Edges Edges bison 43 flex 28 28 26 94 23 gs 94 158 94 105 1478 105 perl 20 VII Intersections of the CMG and CGO popular procedure and edge sets. the CGO one) and the different pruning algorithms used. Although a lot of highly weighted edges maintained their relative positions, the significant performance improvement for CMG came from edges that were promoted higher in the edge list ordering. Table VII shows the intersection between the CMG and CGO popular procedure and edge sets. Columns 2-4 (5- list the CGO and CMG popular procedure (edge) sets along with their intersection. The numbers shown in Table VII are sensitive to the pruning algorithm, but they are compared to better illustrate the differences between the CMG and CGO approach. Although one procedure set is always the superset of the other, the CMG edge list is always larger than the CGO edge list. VI. Acknowledgments We would like to acknowledge the contributions of H. Hashemi and B. Calder to this work. This research was supported by the National Science Foundation CAREER Award Program, by IBM Research and by Microsoft Re-search VII. Conclusions The performance of cache-based memory systems is critical in today's processors. Research has shown that compiler optimizations can significantly reduce memory la- tency, and every opportunity should be taken by the compiler to do so. In this paper we presented two profile-guided algorithms for procedure reordering which take into consideration not only the procedure size but the cache organization as well. While CGO attempts to minimize first-generation conflicts, CMG targets higher generation misses. Both approaches use pruned graph models to guide procedure placement via cache line coloring. The CMG algorithm improved instruction cache miss rates on average by 30% over a static depth first search of procedures, and by 21% over CGO. We also introduced three new sequences (IRPG, IRILG and IRALS) based on the IRG model, to better characterize the contents of each graph model. --R "Evaluating associativity in cpu caches," "Efficient Procedure Mapping using Cache Line Coloring," "Profile-Guided Code Positioning," "An Inter-Reference Gap Model for Temporal Locality in Program Behavior," "Procedure Placement using Temporal Ordering Information," "Program Optimization for Instruction Caches," "Code Reorganization for Instruction Caches," "Procedure Mapping using Static Call Graph Estimation," "Achieving High Instruction Cache Performance with an Optimizing Compiler," "Optimizing instruction cache performance for operating system intensive workloads," "Compile-Time Instruction Cache Optimizations," "Predicting program behavior using real or estimated profiles," "Register allocation and spilling via graph coloring," "Temporal-based Procedure Reordering for improved Instruction Cache Performance," --TR --CTR Altman , David Kaeli , Yaron Sheffer, Guest Editors' Introduction: Welcome to the Opportunities of Binary Translation, Computer, v.33 n.3, p.40-45, March 2000 S. Bartolini , C. A. Prete, Optimizing instruction cache performance of embedded systems, ACM Transactions on Embedded Computing Systems (TECS), v.4 n.4, p.934-965, November 2005 S. Bartolini , C. A. Prete, A proposal for input-sensitivity analysis of profile-driven optimizations on embedded applications, ACM SIGARCH Computer Architecture News, v.32 n.3, p.70-77, June 2004 Mohsen Sharifi , Behrouz Zolfaghari, YAARC: yet another approach to further reducing the rate of conflict misses, The Journal of Supercomputing, v.44 n.1, p.24-40, April 2008

temporal locality;instruction caches;conflict misses;program reordering;graph pruning;graph coloring

297715

Randomized Cache Placement for Eliminating Conflicts.

AbstractApplications with regular patterns of memory access can experience high levels of cache conflict misses. In shared-memory multiprocessors conflict misses can be increased significantly by the data transpositions required for parallelization. Techniques such as blocking which are introduced within a single thread to improve locality, can result in yet more conflict misses. The tension between minimizing cache conflicts and the other transformations needed for efficientparallelization leads to complex optimization problems for parallelizing compilers. This paper shows how the introduction of a pseudorandom element into the cache index function can effectively eliminate repetitive conflict misses and produce a cache where miss ratio depends solely on working set behavior. We examine the impact of pseudorandom cache indexing on processor cycle times and present practical solutions to some of the major implementation issues for this type of cache. Our conclusions are supported by simulations of a superscalar out-of-order processor executing the SPEC95 benchmarks, as well as from cache simulations of individual loop kernels to illustrate specific effects. We present measurements of Instructions committed Per Cycle (IPC) when comparing the performance of different cache architectures on whole-program benchmarks such as the SPEC95 suite.

Introduction If the upward trend in processor clock frequencies during the last ten years is extrapolated over the next ten years, we will see clock frequencies increase by a factor of twenty during that period [1]. However, based on the current 7% per annum reduction in DRAM access times [2], memory latency can be expected to reduce by only 50% in the next ten years. This potential ten-fold increase in the distance to main memory has serious implications for the design of future cache-based memory hierarchies as well as for the architecture of memory devices themselves. Each block of main memory can be placed in exactly one set of blocks in cache. The chosen set is determined by the indexing function. Conventional caches typically extract a field of m bits from the address and use this to select one block from a set of 2 m . Whilst easy to imple- ment, this indexing function is not robust. The principal weakness is its susceptibility to repetitive conflict misses. For example, if C is the number of cache sets and B is the block size, then addresses A 1 and A 2 map to the same cache set if j A 1 =B j C =j A 2 =B j C . If A 1 and A 2 collide on the same cache set, then addresses A 1 A 2 +k also collide in cache, for any integer k , except when There are two common cases when this happens. Firstly, when accessing a stream of addresses fA collides with A i+k , then there may be up to m\Gammak conflict misses in this stream. Secondly, when accessing elements of two distinct arrays b 0 collides with b 1 [j], then b 0 [i+k] collides with under the conditions outlined above. Set-associativity can help to alleviate such conflicts, but is not an effective solution for repetitive and regular conflicts. One of the best ways to control locality in dense matrix computations with large data structures is to use a tiled (or algorithm. This is effectively a re-ordering of the iteration space which increases temporal locality. How- ever, previous work has shown that the conflicts introduced by tiling can be a serious problem [3]. In practice, until now, this has meant that compilers which tile loop nests really ought to compute the maximal conflict-free tile size for given values of B, major array dimension N and cache capacity C. Often this will be too small to make it worth-while tiling a loop, or perhaps the value of N will not be known at compile time. Gosh et al. [4] present a framework for analyzing cache misses in perfectly-nested loops with affine references. They develop a generic technique for determining optimum tile sizes, and methods for determining array padding sizes to avoid conflicts. These methods require solutions to sets of linear Diophantine equations and depend upon there being sufficient information at compile time to find such solutions. Table I highlights the problem of conflict misses with reference to the SPEC95 benchmarks. The programs were compiled with the maximum optimization level and instrumented with the atom tool [5]. A data cache similar to the first-level cache of the Alpha 21164 microprocessor was simulated: 8 KB capacity, 32-byte lines, write-through and no write allocate. For each benchmark we simulated the first 2 operations. Because of the no-write-allocate feature, the tables below refer only to load operations. Table I shows the miss ratio for the following cache or- ganizations: direct-mapped, two-way associative, column- associative [6], victim cache with four victim lines [7], and two-way skewed-associative [8], [9]. Of these schemes, only the two-way skewed-associative cache uses an unconventional indexing scheme, as proposed by its author. For comparison, the miss ratio of a fully-associative cache is shown in the penultimate column. The miss ratio difference between a direct-mapped cache and that of a fully-associative cache is shown in the right-most column of table I, and represents the direct-mapped conflict miss ratio (CMR) [2]. In the case of hydro2d and apsi some organizations exhibit lower miss ratios than a fully-associative cache, due to sub-optimality of lru replacement in a fully-associative cache for these particular programs. Effectively, the direct-mapped conflict miss ratio represents DM 2W CA VC SA FA CMR tomcatv 53.8 48.1 47.0 26.6 22.1 12.5 41.3 su2cor 11.0 9.1 9.3 9.5 9.6 8.9 2.1 hydro2d 17.6 17.1 17.2 17.0 17.1 17.5 0.1 mgrid 3.8 3.6 4.2 3.7 4.1 3.5 0.3 applu 7.6 6.4 6.5 6.9 6.7 5.9 1.7 turb3d 7.5 6.5 6.4 7.0 5.4 2.8 4.7 apsi 15.5 13.3 13.4 10.7 11.5 12.5 3.0 fpppp 8.5 2.7 2.7 7.5 2.2 1.7 6.8 wave 31.8 31.7 30.7 20.1 16.8 13.9 17.9 go 13.4 8.2 8.6 10.9 7.5 4.8 8.6 gcc 10.6 7.2 7.3 8.6 6.6 5.3 5.3 compress 17.1 15.8 16.3 16.2 14.3 13.0 4.1 li 8.6 5.4 5.5 7.2 4.9 3.8 4.8 ijpeg 4.1 3.3 3.1 2.3 1.9 1.2 2.9 perl 10.7 7.3 7.5 9.3 6.9 5.2 5.5 vortex 5.3 2.7 2.7 3.8 1.8 1.4 3.9 Ave Ave (Int) 9.22 6.44 6.67 7.67 5.66 4.42 4.80 Average 15.9 13.8 13.6 11.3 8.66 6.80 9.14 I Cache miss ratios for direct-mapped (DM), 2-way set-associative (2W), column-associative (CA), victim cache (VC), 2-way skewed associative (SA), and fully-associative organizations. Conflict miss ratio (CMR) is also shown. the target reduction in miss ratio that we hope to achieve through improved indexing schemes. The other type of misses, compulsory and capacity, will remain unchanged by the use of randomized indexing schemes. As expected, the improvement of a 2-way set-associative cache over a direct-mapped cache is rather low. The column-associative cache provides a miss ratio similar to that of a two-way set-associative cache. Since the former has a lower access time but requires two cache probes to satisfy some hits, any choice between these two organizations should take into account implementation parameters such as access time and miss penalty. The victim cache removes many conflict misses and outperforms a four-way set-associative cache. Finally, the two-way skewed- associative cache offers the lowest miss ratio. Previous work has shown that it can be significantly more effective than a four-way conventionally-indexed set-associative cache [10]. In this paper we investigate the use of alternative index functions for reducing conflicts and discuss some practical implementation issues. Section II introduces the alternative index functions, and section III evaluates their conflict avoidance properties. In section IV we discuss a number of implementation issues, such as the effect of novel indexing functions on cache access time. Then, in section V, we evaluate the impact of the proposed indexing scheme on the performance of a dynamically-scheduled processor. Finally, in section VI, we draw conclusions from this study. II. Alternative indexing functions The aim of this paper is to show how alternative cache organizations can eliminate repetitive conflict misses. This is analogous to the problem of finding an efficient hashing function. For large secondary or tertiary caches it may be possible to use the virtual address mapping to adjust the location of pages in cache, as suggested by Bershad et al. [11], thus avoiding conflicts dynamically. However, for small first-level caches this effect can only be achieved by using an alternative cache index function. In the field of interleaved memories it is well known that bank conflicts can be reduced by using bank selection functions other than the simple modulo-power-of-two. Lawrie and Vora proposed a scheme using prime-modulus functions [12], Harper and Jump [13], and Sohi [14] proposed skewing functions. The use of xor functions in parallel memory systems was proposed by Frailong et al. [15], and other pseudo-random functions were proposed by Raghavan and Hayes [16], and Rau et al. [17], [18]. These schemes each yield a more or less uniform distribution of requests to banks, with varying degrees of theoretical predictability and implementation cost. In principle each of these schemes could be used to construct a conflict-resistant cache by using them as the indexing function. However, in cache architectures two factors are critical; firstly, the chosen indexing function must have a logically simple im- plementation, and secondly we would like to be able to guarantee good behavior on all regular address patterns - even those that are pathological under a conventional index function. In the commercial domain, the IBM 3033 [19] and the Amdahl 470 [20] made use of xor-mapping functions in order to index the TLB. The first generation HP Precision Architecture processors [21] also used a similar technique. The use of pseudo-random cache indexing has been suggested by other authors. For example, Smith [22] compared a pseudo-random placement against a set-associative place- ment. He concluded that random indexing had a small advantage in most cases, but that the advantages were not significant. In this paper we show that for certain workloads and cache organizations, the advantages can be very large. Hashing the process id with the address bits in order to index the cache was evaluated in a multiprogrammed environment by Agarwal in [23]. Results showed that this scheme could reduce the miss ratio. Perhaps the most well-known alternative cache indexing scheme is the class of bitwise exclusive-OR functions proposed for the skewed associative cache [8]. The bitwise xor mapping computes each bit of the cache index as either one bit of the address or the xor of two bits. Where two such mappings are required different groups of bits are chosen for xor-ing in each case. A two-way skewed-associative cache consists of two banks of the same size that are accessed simultaneously with two different hashing functions. Not only does the associativity help to reduce conflicts but the skewed indexing functions help to prevent repetitive conflicts from occurring. The polynomial modulus function was first applied to cache indexing in [10]. It is best described by first considering the unsigned integer address A in terms of its binary representation This is interpreted as the polynomial defined over the field GF(2). The binary representation of the m-bit cache index R is similarly defined by the GF(2) polynomial R(x) of order less than m such that Effectively R(x) is is an irreducible polynomial of order m and P (x) is such that x i mod P (x) generates all polynomials of order lower than m. The polynomials that fulfil the previous requirements are called Ipoly poly- nomials. Rau showed how the computation of R(x) can be accomplished by the vector-matrix product of the address and an n \Theta m matrix H of single-bit coefficients derived from P (x) [18]. In GF(2), this product is computed by a network of and and xor gates, and if the H-matrix is constant the and gates can be omitted and the mapping then requires just m xor gates with fan-in from 2 to n. In practice we may reduce the number of input address bits to the polynomial mapping function by ignoring some of the upper bits in A. This does not seriously degrade the quality of the mapping function. Ipoly mapping functions have been studied previously in the context of stride-insensitive interleaved memories (see [17], [18]), and have certain provable characteristics of significant value for cache indexing. In [24] it was demonstrated that a skewed Ipoly cache indexing scheme shows a higher degree of conflict resistance than that exhibited by conventional set-associativity or other (non-Ipoly) xor-based mapping functions. Overall, the skewed-associative cache using Ipoly mapping and a pure lru replacement policy achieved a miss ratio within 1% of that achieved by a fully-associative cache. Given the advantage of an Ipoly function over the bitwise xor function, all results presented in this paper use the Ipoly indexing scheme. III. Evaluation of Conflict Resistance The performance of both the integer and floating-point SPEC95 programs has been evaluated for column- associative, two-way set-associative (2W) and two-way skewed-associative organizations using Ipoly indexing functions. In all cases a single-level cache is assumed. The miss ratios of these configurations are shown in table II. Given a conventional indexing function, the direct-mapped (DM) and fully-associative organizations display respectively the lowest and the highest degrees of conflict-resistance of all possible cache architectures. As such they define the bounds within which novel indexing schemes should be evaluated. Their miss ratios are shown in the right-most two columns of table II. The column-associative cache has access-time characteristics similar to a direct-mapped cache but has some degree of pseudo-associativity - each address can map to one of Ipoly indexing mod 2 k col. assoc. 2-way skewed indexing spl lru 2W plru lru FA DM su2cor 10.5 9.1 9.9 9.4 9.4 8.9 11.0 hydro2d 17.6 17.2 17.1 17.0 17.1 17.5 17.6 mgrid 5.1 4.2 3.8 4.5 4.1 3.5 3.8 applu 7.3 6.5 6.9 6.8 6.4 5.9 7.6 turb3d 8.1 6.0 4.8 4.5 4.2 2.8 7.5 apsi 12.2 11.2 11.4 11.0 10.6 12.5 15.5 fpppp 4.0 2.7 2.8 2.1 2.3 1.7 8.5 wave ? 14.6 13.8 14.2 13.9 13.7 13.9 31.8 go 9.6 6.6 8.6 7.5 6.7 4.8 13.4 gcc 8.2 6.3 7.2 6.7 6.1 5.3 10.6 compress 14.5 13.5 13.7 13.9 13.4 13.0 17.1 li 5.5 4.5 6.1 4.9 4.5 3.8 8.6 ijpeg 1.8 1.3 1.7 1.5 1.4 1.2 4.1 perl 8.5 6.7 8.8 7.1 6.4 5.2 10.7 vortex 2.7 1.7 2.0 1.8 1.6 1.4 5.3 Ave Ave (Int) 6.68 5.22 6.26 5.55 5.09 4.42 9.22 Ave ? 13.2 11.4 12.3 11.6 11.3 11.4 47.3 Average 8.77 7.39 7.99 7.47 7.14 6.80 15.9 II Miss ratios for Ipoly indexing on SPEC95 benchmarks. two locations in the cache, but initially only one is probed. The column labelled spl represents a cache which swaps data between the two locations to increase the percentage of a hit on the first probe. It also uses a realistic pseudo-lru replacement policy. The cache reported in the column labelled lru does not swap data between columns and uses an unrealistic pure lru replacement policy [10]. It is to be expected that a two-way set-associative cache will be capable of eliminating many random conflicts. How- ever, a conventionally-indexed set-associative cache is not able to eliminate pathological conflict behavior as it has limited associativity and a naive indexing function. The performance of a two-way set-associative cache can be improved by simply replacing the index function, whilst retaining all other characteristics. Conventional lru replacement can still be used, as the indexing function has no impact on replacement for this cache organization. For two programs the two-way Ipoly cache has a lower miss ratio than a fully-associative cache. This is again due to the sub-optimality of lru replacement in the fully-associative cache, and is a common anomaly in programs with negligible conflict misses. The final cache organization shown in table II is the two-way skewed-associative cache proposed originally by Seznec [8]. In its original form it used two bitwise xorindexing functions. Our version uses Ipoly indexing functions, as proposed in [10] and [24]. In this case two distinct Ipoly functions are used to construct two distinct cache indices from each address. Pure lru is difficult to implement in a skewed-associative cache, so here we present results for an cache which uses a realistic pseudo-lru policy (labelled plru) and a cache which uses an unrealistic pure lru policy (labelled lru). This organization produces the lowest conflict miss ratio, down from 4.8% to 0.67% for SPECint, and from 12.61% to 0.07% for SPECfp. It is striking that the performance improvement is dominated by three programs (tomcatv, swim and wave). These effectively exhibit pathological conflict miss ratios under conventional indexing schemes. Studies by Olukotun et al. [25], have shown that the data cache miss ratio in tomcatv wastes 56% and 40% of available IPC in 6-way and 2-way superscalar processors respectively. Tiling will often introduce extra cache conflicts, the elimination of which is not always possible through software. Now that we have alternative indexing functions that exhibit conflict avoidance properties we can use these to avoid these induced conflicts. The effectiveness of Ipoly indexing for tiled loops was evaluated by simulating the cache behavior of a variety of tiled loop kernels. Here we present a small sample of results to illustrate the general outcome. Figures show the miss ratios observed in two tiled matrix multiplication kernels where the original matrices were square and of dimensions 171 and 256 respectively. Tile sizes were varied from 2 \Theta 2 up to 16 \Theta 16 to show the effect of conflicts occurring in caches that are direct-mapped (a1), 2-way set-associative (a2), fully-associative (fa) and skewed 2-way Ipoly (Hp-Sk). The tiled working set divided by cache capacity measures the fraction of the cache occupied by a single tile. Cache capacity is 8 KBytes, with 32-byte lines. For dimension 171 the miss ratio initially falls for all caches as tile size increases. This is due to increasing spatial locality, up to the point where self conflicts begin to occur in the conventionally-indexed direct-mapped and two-way set-associative caches. The fully-associative cache suffers no self-conflicts and its miss ratio decreases monotonically to less than 1% at 50% loading. The behavior of the skewed 2-way Ipoly cache tracks the fully-associative cache closely. The qualitative difference between the Ipoly cache and a conventional two-way cache is clearly visible. For dimension 256 the product array and the multiplicand array are positioned in memory so that cross-conflicts occur in addition to self-conflicts. Hence the direct-mapped and 2-way set associative caches experience little spatial locality. However, the Ipoly cache is able to eliminate cross-conflicts as well as self-conflicts, and it again tracks the fully-associative cache. IV. Implementation Issues The logic of the GF(2) polynomial modulus operation presented in section II defines a class of hash functions which compute the cache placement of an address by combining subsets of the address bits using xor gates. This means that, for example, bit 0 of the cache index may be Working Set / Capacity Miss ratio 0% 10% 20% 30% 40% 50% Hp-Sk Fig. 1. Miss ratio versus cache loading for 171 \Theta 171 matrix multiply. Working Set / Capacity Miss ratio 0% 10% 20% 30% 40% 50% Fig. 2. Miss ratio versus cache loading for 256 \Theta 256 matrix multiply. computed as the xor of bits 0, 11, 14, and 19 of the original address. The choice of polynomial determines which bits are included in each set. The implementation of such a function for a cache with an 8-bit index would require just eight xor gates with fan-in of 3 or 4. Whilst this appears remarkably simple, there is more to consider than just the placement function. Firstly, the function itself uses address bits beyond the normal limit imposed by typical minimum page size restriction. Sec- ondly, the use of pseudo-random placement in a multi-level memory hierarchy has implications for the maintenance of Inclusion. In [24] we explain these two issues in more depth, and show how the virtual-real two-level cache hierarchy proposed by Wang et al. [26] provides a clean solution to both problems. A cache memory access in a conventional organization normally computes its effective address by adding two reg- isters, or a register plus a displacement. Ipoly indexing implies additional circuitry to compute the index from the effective address. This circuitry consists of several xor gates that operate in parallel and therefore the total delay is just the delay of one gate. Each xor gate has a number of inputs that depend on the particular polynomial being used. For the experiments reported in this paper the number of inputs is never higher than 5. The xor gating required by the Ipoly mapping may increase the critical path length within the processor pipeline. However, any delay will be short since all bits of the index can be computed in parallel. Moreover, we show later that even if this additional delay induces a full cycle penalty in the cache access time, the Ipoly mapping provides a significant overall performance improvement. Memory address prediction can be also used to avoid the penalty introduced by the xor delay when it lengthens the critical path. Memory addresses have been shown to be highly predictable. For instance, in [27] it was shown that the addresses of about 75% of the dynamically executed memory instructions from the SPEC95 suite can be predicted with a simple tabular scheme which tracks the last address produced by a given instruction and its last stride. A similar scheme, that could be used to give an early prediction of the line that is likely to be accessed by a given load instruction, is outlined below. The processor incorporates a table indexed by the instruction address. Each entry stores the last address and the predicted stride for some recently executed load in- struction. In the fetch stage, this table is accessed with the program counter. In the decode stage, the predicted address is computed and the xor functions are performed to compute the predicted cache line. This can be done in one cycle since the xor can be performed in parallel with the computation of the most-significant bits of the effeec- tive address. When the instruction is subsequently issued to the memory unit it uses the predicted line number to access the cache in parallel with the actual address and line computation. If the predicted line turns out to be incor- rect, the cache access is repeated with the actual address. Otherwise, the data provided by the speculative access can be loaded into the destination register. A number of previous papers have suggested address prediction as a means to reduce memory latency [28], [29], [30], or to execute memory instructions and their dependent instructions speculatively [31], [27], [32]. In the case of a miss-speculation, a recovery mechanism similar to that used by branch prediction schemes is then used to squash miss-speculated instructions. V. Effect of Ipoly indexing on IPC In order to verify the impact of polynomial mapping on realistic microprocessor architectures we have developed a parametric simulator for a four-way superscalar processor with out-of-order execution. Table III summarizes the functional units and their latencies used in these experiments. The reorder buffer contained 32 entries, and there were two separate physical register files (FP and In- teger), each with 64 physical registers. The processor had a lockup-free data cache [33] that allowed 8 outstanding misses to different cache lines. Cache capacities of 8 KB and were simulated with 2-way associativity and 32-byte lines. The cache was write-through and no-write- allocate. The cache had two ports, each with a two-cycle time and a miss penalty of 20 cycles. This was connected by a 64-bit data bus to an infinite level-two cache. Data dependencies through memory were speculated using a mechanism similar to the arb of the Multiscalar [34] and the HP PA-8000 [35]. A branch history table with 2K entries and 2-bit saturating counters was used for branch prediction. Functional unit Latency Repeat rate simple FP 4 1 III Functional units and instruction latencies The memory address prediction scheme was implemented by a direct-mapped table with 1K entries, indexed by instruction address. To reduce cost the entries were not tagged, although this increases interference in the ta- ble. Each entry contained the last effective address of the most recent load instruction to index into that table entry, together with the last observed stride. In addition, each entry contained a 2-bit saturating counter to assign confidence to the prediction. Only when the most-significant bit of the counter is set would the prediction be considered correct. The address field was updated for each new reference regardless of the prediction. However, the stride field was updated only when the counter went below after two consecutive mispredictions. Table IV shows the IPC and miss ratios for six configurations 1 . All IPC averages are computed using an equally-weighted harmonic mean. The baseline configuration is an 8 KB cache with conventional indexing and no address prediction (np, 3rd column). The average IPC for this configuration is 1.27 from an average miss ratio of 16.53%. With Ipoly indexing the average miss ratio falls to 9.68%. If the xor gates are not in the critical path IPC rises to 1.33 (nx, 5th column). Conversely, if the xor gates are in the critical path, and a one cycle penalty in the cache access time is assumed, the resulting IPC is 1.29 (wx, 6th column). How- ever, if memory address prediction is then introduced (wp, 7th columnn) IPC is the same as for a cache without the xor gates in the critical path (nx). Hence, the memory address prediction scheme can offset the penalty introduced by the additional delay of the xor gates when they are in the critical path, even under the conservative assumption that whole cycle of latency is added to each load instruc- tion. Finally, table IV also shows the performance of a set-associative cache without Ipoly indexing 1 For each configuration we simulated 10 8 instructions after skipping the first 2 \Theta 10 9 . (2nd column). Notice that the addition of Ipoly indexing to an 8 KB cache yields over 60% of the IPC increase that can be obtained by doubling the cache size. indexing Ipoly indexing su2cor y 1.28 1.24 1.26 1.24 1.21 1.25 hydro2d y 1.14 1.13 1.15 1.13 1.11 1.15 mgrid y 1.63 1.61 1.63 1.57 1.55 1.59 applu y 1.51 1.50 1.53 1.50 1.46 1.52 turb3d y 1.85 1.80 1.82 1.81 1.78 1.82 apsi y 1.13 1.08 1.09 1.08 1.07 1.09 fpppp y 2.14 2.00 2.00 1.98 1.93 1.94 wave ? 1.37 1.26 1.28 1.51 1.48 1.54 go y 1.00 0.87 0.88 0.87 0.83 0.84 compress y 1.13 1.12 1.13 1.11 1.07 1.10 li y 1.40 1.30 1.32 1.33 1.26 1.31 ijpeg y 1.31 1.28 1.28 1.29 1.28 1.30 perl y 1.45 1.26 1.27 1.24 1.19 1.21 vortex y 1.39 1.27 1.28 1.30 1.25 1.27 Ave Ave (Int) 1.29 1.19 1.20 1.20 1.15 1.17 Ave ? 1.28 1.11 1.13 1.46 1.42 1.49 Ave y 1.38 1.30 1.32 1.30 1.27 1.30 Average 1.36 1.27 1.28 1.33 1.29 1.33 IV Comparative IPC measurements (simulated). These IPC measurements exhibit small absolute differ- ences, but this is because the benefit of Ipoly indexing is perceived by a only small subset of the benchmark pro- grams. Most programs in SPEC95 exhibit low conflict miss ratios. In fact the SPEC95 conflict miss ratio for an 8 KB 2-way set-associative cache is less than 4% for all programs except tomcatv, swim and wave5. The two penultimate rows of table IV show independent IPC averages for the benchmarks with high conflict miss ratios (Ave ?), and those with low conflict miss ratios (Ave y). This highlights the ability of polynomial mapping to reduce the miss ratio and significantly boost the performance of problem cases. One can see that the polynomial mapping provides a significant 27% improvement in IPC for the three bad programs even if the xor gates are in the critical path and memory address prediction is not used. With memory address prediction Ipoly indexing yields an IPC improvement of 33% compared with that of a conventional cache of the same capacity, and 16% higher than that of a conventional cache with twice the capacity. Notice that the polynomial mapping scheme with prediction is even better than the organization without prediction where the xor gates do not extend the critical path. This is due to the fact that the memory address prediction scheme reduces by one cycle the effective cache hit time when the predictions are correct, since the address computation is overlapped with the cache access (the computed address is used to verify that the prediction was correct). However, the main benefits observed come from the reduction in conflict misses. To isolate the different effects we have also simulated an organization with the memory address prediction scheme and conventional indexing for an 8 KB cache (wp, column 4). If we compare the IPC of this organization with that in column 3, we see that the benefits of the memory address prediction scheme due solely to the reduction of the hit time are almost negligible. This confirms that the improvement observed in the Ipoly indexing scheme with address prediction derives from the reduction in conflict misses. The averages for the fifteen programs which exhibit low levels of conflict misses show a small (1.7%) deterioration in average IPC when Ipoly indexing is used and the xor gates are in the critical path. This is due to a slight increase in the average hit time rather than an overall increase in miss ratio (which on average falls by 2%). For these programs the reduction in aggregated miss penalty does not outweigh the slight extension in critical path length. VI. Conclusions In this paper we have discussed the problem of cache conflict misses and surveyed the options for reducing or eliminating those conflicts. We have described pseudo-random indexing schemes based on polynomial modulus functions, and have shown them to be robust enough to virtually eliminate the repetitive cache conflicts caused by bad strides inherent in some SPEC95 benchmarks, as well as eliminating those introduced into an application by the tiling of loop nests. We have highlighted the major implementation issues that arise from the use of such novel indexing schemes. For example, Ipoly indexing uses more address bits than a conventional cache to compute the cache index. Also, the use of different indexing functions at level-1 and level-2 caches results in the occasional eviction at level-1 simply to maintain Inclusion. We have explained that both of these problems can be solved using a two-level virtual-real cache hierarchy. Finally, we have proposed a memory address prediction scheme to avoid the penalty due to the small potential delay in the critical path introduced by the pseudo-random indexing function. Detailed simulations of an out-of-order superscalar processor have demonstrated that programs with significant numbers of conflict misses in a conventional 8 KB 2-way skewed-associative cache perceive IPC improvements of 33% (with address prediction) or 27% (without address prediction). This is up to 16% higher than the IPC improvements obtained simply by doubling the cache capac- ity. Furthermore, from the programs we analyzed, those that do not experience significant conflict misses on average see only a 1.7% reduction in IPC when Ipoly indexing appears on the critical path for computing the effective ad- dress, and address prediction is used. If the indexing logic does not appear on the critical path no deterioration in overall average performance is experienced by those programs We believe the key contribution of pseudo-random indexing is the resulting predictability of cache behavior. In our experiments we found that Ipoly indexing reduces the standard deviation of miss ratios across SPEC95 from 18.49 to 5.16. This could be beneficial in real-time systems where unpredictable timing, caused by the possibility of pathological miss ratios, presents problems. If conflict misses are eliminated, the miss ratio depends solely on compulsory and capacity misses, which in general are easier to predict and control. Conflict avoidance could also be beneficial when iteration-space tiling is used to improve data locality. VII. Acknowledgments This work was supported in part by the European Commission esprit project 24942, by the British Council (grant 1016), and by the Spanish Ministry of Education (Acci'on Integrada Hispano-Brit'anica 202 CYCIT TIC98- 0511). The authors would like to express their thanks to Jose Gonz'alez and Joan Manuel Parcerisa for their help with the simulation software, and to the anonymous referees for their helpful comments. --R "The national technology roadmap for semiconductors," Computer Architecture: A Quantitative Approach. "The cache performance and optimization of blocked algorithms," "Cache miss equations: An analytic representation of cache misses," "ATOM: A system for building customized program analysis tools," "Column-associative caches: A technique for reducing the miss rate of direct-mapped caches," "Improving direct-mapped cache performance by the addition of a small fully-associative cache and prefetch buffers," "A case for two-way skewed associative caches," "Skewed-associative caches," "Elimi- nating cache conflict misses through xor-based placement func- tions," "Avoiding cache conflict misses dynamically in large direct-mapped caches," "The prime memory system for array access," "Vector access performance in parallel memories," "Logical data skewing schemes for interleaved memories in vector processors," "xor-schemes: A flexible data organization in parallel memories," "On randomly interleaved memo- ries," "The Cydra 5 stride- insensitive memory system," "Pseudo-randomly interleaved memories," Amdahl Corp. "Hardware design of the first HP Precision Architecture computers," "Cache memories," Analysis of Cache Performance for Operating Systems and Multiprogramming "The design and performance of a conflict-avoiding cache," "The case for a single-chip multiprocessor," "Organization and performance of a two-level virtual-real cache hierarchy," "Speculative execution via address prediction and data prefetching," "Hardware support for hiding cache latency," "Streamlining data cache access with fast address calculation," "Zero-cycle loads: Microarchitecture support for reducing load latency," "Memory address prediction for data speculation," "The performance potential of data dependence speculation and collapsing," "Lockup-free instruction fetch/prefetch cache organi- zation," "ARB: A mechanism for dynamic reordering of memory references," "Advanced performance features of the 64-bit PA- 8000," --TR --CTR S. Bartolini , C. A. Prete, A proposal for input-sensitivity analysis of profile-driven optimizations on embedded applications, ACM SIGARCH Computer Architecture News, v.32 n.3, p.70-77, June 2004 K. Patel , E. Macii , L. Benini , M. Poncino, Reducing cache misses by application-specific re-configurable indexing, Proceedings of the 2004 IEEE/ACM International conference on Computer-aided design, p.125-130, November 07-11, 2004 Hans Vandierendonck , Philippe Manet , Jean-Didier Legat, Application-specific reconfigurable XOR-indexing to eliminate cache conflict misses, Proceedings of the conference on Design, automation and test in Europe: Proceedings, March 06-10, 2006, Munich, Germany Hans Vandierendonck , Koen De Bosschere, XOR-Based Hash Functions, IEEE Transactions on Computers, v.54 n.7, p.800-812, July 2005 Mathias Spjuth , Martin Karlsson , Erik Hagersten, Skewed caches from a low-power perspective, Proceedings of the 2nd conference on Computing frontiers, May 04-06, 2005, Ischia, Italy Wang , Nelson Passos, Improving cache hit ratio by extended referencing cache lines, Journal of Computing Sciences in Colleges, v.18 n.4, p.118-123, April G. E. Suh , L. Rudolph , S. Devadas, Dynamic Partitioning of Shared Cache Memory, The Journal of Supercomputing, v.28 n.1, p.7-26, April 2004 Hans Vandierendonck , Koen De Bosschere, Highly accurate and efficient evaluation of randomising set index functions, Journal of Systems Architecture: the EUROMICRO Journal, v.48 n.13-15, p.429-452, May G. Edward Suh , Srinivas Devadas , Larry Rudolph, Analytical cache models with applications to cache partitioning, Proceedings of the 15th international conference on Supercomputing, p.1-12, June 2001, Sorrento, Italy Rui Min , Yiming Hu, Improving Performance of Large Physically Indexed Caches by Decoupling Memory Addresses from Cache Addresses, IEEE Transactions on Computers, v.50 n.11, p.1191-1201, November 2001 S. Bartolini , C. A. Prete, Optimizing instruction cache performance of embedded systems, ACM Transactions on Embedded Computing Systems (TECS), v.4 n.4, p.934-965, November 2005

cache architectures;performance evaluation;conflict avoidance

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The Impact of Exploiting Instruction-Level Parallelism on Shared-Memory Multiprocessors.

AbstractCurrent microprocessors incorporate techniques to aggressively exploit instruction-level parallelism (ILP). This paper evaluates the impact of such processors on the performance of shared-memory multiprocessors, both without and with the latency-hiding optimization of software prefetching. Our results show that, while ILP techniques substantially reduce CPU time in multiprocessors, they are less effective in removing memory stall time. Consequently, despite the inherent latency tolerance features of ILP processors, we find memory system performance to be a larger bottleneck and parallel efficiencies to be generally poorer in ILP-based multiprocessors than in previous generation multiprocessors. The main reasons for these deficiencies are insufficient opportunities in the applications to overlap multiple load misses and increased contention for resources in the system. We also find that software prefetching does not change the memory bound nature of most of our applications on our ILP multiprocessor, mainly due to a large number of late prefetches and resource contention. Our results suggest the need for additional latency hiding or reducing techniques for ILP systems, such as software clustering of load misses and producer-initiated communication.

Introduction Shared-memory multiprocessors built from commodity microprocessors are being increasingly used to provide high performance for a variety of scientific and commercial ap- plications. Current commodity microprocessors improve performance by using aggressive techniques to exploit high levels of instruction-level parallelism (ILP). These techniques include multiple instruction issue, out-of-order (dy- namic) scheduling, non-blocking loads, and speculative ex- ecution. We refer to these techniques collectively as ILP techniques and to processors that exploit these techniques as ILP processors. Most previous studies of shared-memory multiproces- This work is supported in part by an IBM Partnership award, Intel Corporation, the National Science Foundation under Grant No. CCR-9410457, CCR-9502500, CDA-9502791, and CDA-9617383, and the Texas Advanced Technology Program under Grant No. 003604-025. Sarita Adve is also supported by an Alfred P. Sloan Research Fellowship, Vijay S. Pai by a Fannie and John Hertz Foundation Fellowship, and Parthasarathy Ranganathan by a Lodieska Stockbridge Vaughan Fellowship. This paper combines results from two previous conference papers [11], [12], using a common set of system parameters, a more aggressive MESI (versus MSI) cache-coherence protocol, a more aggressive compiler (the better of SPARC SC 4.2 and gcc 2.7.2 for each application, rather than gcc 2.5.8), and full simulation of private memory references. sors, however, have assumed a simple processor with single- issue, in-order scheduling, blocking loads, and no specu- lation. A few multiprocessor architecture studies model state-of-the-art ILP processors [2], [7], [8], [9], but do not analyze the impact of ILP techniques. To fully exploit recent advances in uniprocessor technology for shared-memory multiprocessors, a detailed analysis of how ILP techniques affect the performance of such systems and how they interact with previous optimizations for such systems is required. This paper evaluates the impact of exploiting ILP on the performance of shared-memory multiprocessors, both without and with the latency-hiding optimization of software prefetching. For our evaluations, we study five applications using detailed simulation, described in Section II. Section III analyzes the impact of ILP techniques on the performance of shared-memory multiprocessors without the use of software prefetching. All our applications see performance improvements from the use of current ILP techniques, but the improvements vary widely. In partic- ular, ILP techniques successfully and consistently reduce the CPU component of execution time, but their impact on the memory stall time is lower and more application- dependent. Consequently, despite the inherent latency tolerance features integrated within ILP processors, we find memory system performance to be a larger bottleneck and parallel efficiencies to be generally poorer in ILP-based multiprocessors than in previous-generation multiproces- sors. These deficiencies are caused by insufficient opportunities in the application to overlap multiple load misses and increased contention for system resources from more frequent memory accesses. Software-controlled non-binding prefetching has been shown to be an effective technique for hiding memory latency in simple processor-based shared memory systems [6]. Section IV analyzes the interaction between software prefetching and ILP techniques in shared-memory multiprocessors. We find that, compared to previous-generation systems, increased late prefetches and increased contention for resources cause software prefetching to be less effective in reducing memory stall time in ILP-based systems. Thus, even after adding software prefetching, most of our applications remain largely memory bound on the ILP-based system. our results suggest that, compared to previous-generation 2shared-memory systems, ILP-based systems have a greater need for additional techniques to tolerate or reduce memory latency. Specific techniques motivated by our results include clustering of load misses in the applications to increase opportunities for load misses to overlap with each other, and techniques such as producer-initiated communication that reduce latency to make prefetching more effective (Section V). II. Methodology A. Simulated Architectures To determine the impact of ILP techniques on multiprocessor performance, we compare two systems - ILP and Simple - equivalent in every respect except the processor used. The ILP system uses state-of-the-art ILP processors while the Simple system uses simple processors (Section II- A.1). We compare the ILP and Simple systems not to suggest any architectural tradeoffs, but rather, to understand how aggressive ILP techniques impact multiprocessor per- formance. Therefore, the two systems have identical clock rates, and include identical aggressive memory and network configurations suitable for the ILP system (Section II-A.2). Figure 1 summarizes all the system parameters. A.1 Processor models The ILP system uses state-of-the-art processors that include multiple issue, out-of-order (dynamic) scheduling, non-blocking loads, and speculative execution. The Simple system uses previous-generation simple processors with single issue, in-order (static) scheduling, and blocking loads, and represents commonly studied shared-memory systems. Since we did not have access to a compiler that schedules instructions for our in-order simple processor, we assume single-cycle functional unit latencies (as also assumed by most previous simple-processor based shared-memory stud- ies). Both processor models include support for software-controlled non-binding prefetching to the L1 cache. A.2 Memory Hierarchy and Multiprocessor Configuration We simulate a hardware cache-coherent, non-uniform memory access (CC-NUMA) shared-memory multiprocessor using an invalidation-based, four-state MESI directory coherence protocol [4]. We model release consistency because previous studies have shown that it achieves the best performance [9]. The processing nodes are connected using a two-dimensional mesh network. Each node includes a proces- sor, two levels of caches, a portion of the global shared-memory and directory, and a network interface. A split-transaction bus connects the network interface, directory controller, and the rest of the system node. Both caches use a write-allocate, write-back policy. The cache sizes are chosen commensurate with the input sizes of our ap- plications, following the methodology described by Woo et al. [14]. The primary working sets for our applications fit in the L1 cache, while the secondary working sets do not fit in the L2 cache. Both caches are non-blocking and use miss Processor parameters Clock rate 300 MHz Fetch/decode/retire rate 4 per cycle Instruction window (re- order buffer) sizeMemory queue size Outstanding branches 8 Functional units 2 ALUs, 2 FPUs, 2 address generation units; all 1 cycle latency Memory hierarchy and network parameters MSHRs, 64-byte line L2 cache 64 KB, 4-way associative, 1 port, 8 MSHRs, 64-byte line, pipelined Memory 4-way interleaved, ns access time, Bus 100 MHz, 128 bits, split transaction Network 2D mesh, 150MHz, 64 bits, per hop flit delay of 2 network cycles Nodes in multiprocessor 8 Resulting contentionless latencies (in processor cycles) Local memory 45 cycles Remote memory 140-220 cycles Cache-to-cache transfer 170-270 cycles Fig. 1. System parameters. status holding registers (MSHRs) [3] to store information on outstanding misses and to coalesce multiple requests to the same cache line. All multiprocessor results reported in this paper use a configuration with 8 nodes. B. Simulation Environment We use RSIM, the Rice Simulator for ILP Multipro- cessors, to model the systems studied [10]. RSIM is an execution-driven simulator that models the processor pipelines, memory system, and interconnection network in detail, including contention at all resources. It takes application executables as input. To speed up our simulations, we assume that all instructions hit in the instruction cache. This assumption is reasonable since all our applications have very small instruction footprints. C. Performance Metrics In addition to comparing execution times, we also report the individual components of execution time - CPU, data memory stall, and synchronization stall times - to characterize the performance bottlenecks in our systems. With ILP processors, it is unclear how to assign stall time to specific instructions since each instruction's execution may be overlapped with both preceding and following instructions. We use the following convention, similar to previous work (e.g., [5]), to account for stall cycles. At every cycle, we calculate the ratio of the instructions retired from the instruction window in that cycle to the maximum retire rate of the processor and attribute this fraction of the cycle to the busy time. The remaining fraction of the cycle is attributed as stall time to the first instruction that could not be retired that cycle. We group the busy time and functional unit (non-memory) stall time together as CPU time. Henceforth, we use the term memory stall time to denote the data memory stall component of execution time. In the first part of the study, the key metric used to Application Input Size LU, LUopt 256x256 matrix, block 8 FFT, FFTopt 65536 points Mp3d 50000 particles Water 512 molecules Fig. 2. Applications and input sizes. evaluate the impact of ILP is the ratio of the execution time with the Simple system relative to that achieved by the ILP system, which we call the ILP speedup. For detailed analysis, we analogously define an ILP speedup for each component of execution time. D. Applications Figure 2 lists the applications and the input sets used in this study. Radix, LU, and FFT are from the SPLASH- suite [14], and Water and Mp3d are from the SPLASH suite [13]. These five applications and their input sizes were chosen to ensure reasonable simulation times. (Since RSIM models aggressive ILP processors in detail, it is about 10 times slower than simple-processor-based shared-memory simulators.) LUopt and FFTopt are versions of LU and FFT that include ILP-specific optimizations that can potentially be implemented in a compiler. Specifically, we use function inlining and loop interchange to move load misses closer to each other so that they can be overlapped in the ILP processor. The impact of these optimizations is discussed in Sections III and V. Both versions of LU are also modified slightly to use flags instead of barriers for better load balance. Since a SPARC compiler for our ILP system does not ex- ist, we compiled our applications with the commercial Sun SC 4.2 or the gcc 2.7.2 compiler (based on better simulated ILP system performance) with full optimization turned on. The compilers' deficiencies in addressing the specific instruction grouping rules of our ILP system are partly hidden by the out-of-order scheduling in the ILP processor. 2 III. Impact of ILP Techniques on Performance This section analyzes the impact of ILP techniques on multiprocessor performance by comparing the Simple and ILP systems, without software prefetching. A. Overall Results Figures 3 and 4 illustrate our key overall results. For each application, Figure 3 shows the total execution time and its three components for the Simple and ILP systems (normalized to the total time on the Simple system). Ad- ditionally, at the bottom, the figure also shows the ILP speedup for each application. Figure 4 shows the parallel efficiency 3 of the ILP and Simple systems expressed as a percentage. These figures show three key trends: ffl ILP techniques improve the execution time of all our ap- plications. However, the ILP speedup shows a wide vari- 2 To the best of our knowledge, the key compiler optimization identified in this paper (clustering of load misses) is not implemented in any current superscalar compiler. 3 The parallel efficiency for an application on a system with N processors is defined as Execution time on uniprocessor Execution time on multiprocessor \Theta 1 ation (from 1.29 in Mp3d to 3.54 in LUopt). The average ILP speedup for the original applications (i.e., not including LUopt and FFTopt) is only 2.05. ffl The memory stall component is generally a larger part of the overall execution time in the ILP system than in the Simple system. ffl Parallel efficiency for the ILP system is less than that for the Simple system for all applications. We next investigate the reasons for the above trends. B. Factors Contributing to ILP Speedup Figure 3 indicates that the most important components of execution time are CPU time and memory stalls. Thus, ILP speedup will be shaped primarily by CPU ILP speedup and memory ILP speedup. Figure 5 summarizes these speedups (along with the total ILP speedup). The figure shows that the low and variable ILP speedup for our applications can be attributed largely to insufficient and variable memory ILP speedup; the CPU ILP speedup is similar and significant among all applications (ranging from 2.94 to 3.80). More detailed data shows that for most of our applications, memory stall time is dominated by stalls due to loads that miss in the L1 cache. We therefore focus on the impact of ILP on (L1) load misses below. The load miss ILP speedup is the ratio of the stall time due to load misses in the Simple and ILP systems, and is determined by three factors, described below. The first factor increases the speedup, the second decreases it, while the third may either increase or decrease it. ffl Load miss overlap. Since the Simple system has blocking loads, the entire load miss latency is exposed as stall time. In ILP, load misses can be overlapped with other useful work, reducing stall time and increasing the ILP load miss speedup. The number of instructions behind which a load miss can overlap is, however, limited by the instruction window size; further, load misses have longer latencies than other instructions in the instruction window. There- fore, load miss latency can normally be completely hidden only behind other load misses. Thus, for significant load miss ILP speedup, applications should have multiple load clustered together within the instruction window to enable these load misses to overlap with each other. Contention. Compared to the Simple system, the ILP system can see longer latencies from increased contention due to the higher frequency of misses, thereby negatively affecting load miss ILP speedup. ffl Change in the number of misses. The ILP system may see fewer or more misses than the Simple system because of speculation or reordering of memory ac- cesses, thereby positively or negatively affecting load miss ILP speedup. All of our applications except LU see a similar number of cache misses in both the Simple and ILP case. LU sees 2.5X fewer misses in ILP because of a reordering of accesses that otherwise conflict. When the number of misses does not change, the ILP system sees (? 1) load miss ILP speedup if the load miss overlap exploited by ILP outweighs any additional latency from contention. We illustrate the d execution time ILP Simple ILP41.5 2.41X Simple ILP39.0 2.56X Simple ILP34.0 2.94X Simple ILP28.2 3.54X Simple ILP78.2 1.29X Simple ILP75.3 Simple ILP44.9 2.30X Memory CPU speedup FFT FFTopt LU LUopt Mp3d Radix Water Fig. 3. Impact of ILP on multiprocessor performance. ||||||% Parallel efficiency 86 82 FFT LU Water Simple ILP Fig. 4. Impact of ILP on parallel efficiency. Fig. 5. ILP speedup for total execution time, CPU time, and memory stall time in the multiprocessor system. effects of load miss overlap and contention using the two applications that best characterize them: LUopt and Radix. Figure 6(a) provides the average load miss latencies for LUopt and Radix in the Simple and ILP systems, normalized to the Simple system latency. The latency shown is the total miss latency, measured from address generation to data arrival, including the overlapped part (in ILP) and the exposed part that contributes to stall time. The difference in the bar lengths of Simple and ILP indicates the additional latency added due to contention in ILP. Both of these applications see a significant latency increase from resource contention in ILP. However, LUopt can overlap all its additional latency, as well as a large portion of the base (Simple) latency, thus leading to a high memory ILP speedup. On the other hand, Radix cannot overlap its additional latency; thus, it sees a load miss slowdown in the ILP configuration. We use the data in Figures 6(b) and (c) to further investigate the causes for the load miss overlap and contention- related latencies in these applications. Causes for load miss overlap. Figure 6(b) shows the ILP system's L1 MSHR occupancy due to load misses for LUopt and Radix. Each curve shows the fraction of total time for which at least N MSHRs are occupied by load misses, for each possible N (on the X axis). This figure shows that LUopt achieves significant overlap of load misses, with up to 8 load miss requests outstanding simultaneously at various times. In contrast, Radix almost never has more than 1 outstanding load miss at any time. This difference arises because load misses are clustered together in the instruction window in LUopt, but typically separated by too many instructions in Radix. Causes for contention. Figure 6(c) extends the data of Figure 6(b) by displaying the total MSHR occupancy for both load and store misses. The figure indicates that Radix has a large amount of store miss overlap. This overlap does not contribute to an increase in memory ILP speedup since store latencies are already hidden in both the Simple and ILP systems due to release consistency. The store miss overlap, however, increases contention in the memory hi- erarchy, resulting in the ILP memory slowdown in Radix. In LUopt, the contention-related latency comes primarily from load misses, but its effect is mitigated since overlapped load misses contribute to reducing memory stall time. C. Memory stall component and parallel efficiency Using the above analysis, we can see why the ILP system generally sees a larger relative memory stall time component Figure and a generally poorer parallel efficiency Figure than the Simple system. Since memory ILP speedup is generally less than CPU ILP speedup, the memory component becomes a greater fraction of total execution time in the ILP system than in the Simple system. To understand the reduced parallel effi- ciency, Figure 7 provides the ILP speedups for the uniprocessor configuration for reference. The uniprocessor also generally sees lower memory ILP speedups than CPU ILP speedups. However, the impact of the lower memory ILP speedup is higher in the multiprocessor because the longer latencies of remote misses and increased contention result in a larger relative memory component in the execution time (relative to the uniprocessor). Additionally, the dichotomy between local and remote miss latencies in a multiprocessor often tends to decrease memory ILP speedup (relative to the uniprocessor), because load misses must be overlapped not only with other load misses but with load misses with similar latencies 4 . Thus, overall, the multiprocessor system is less able to exploit ILP features than the corresponding uniprocessor system for most applications. 4 FFT and FFTopt see better memory ILP speedups in the multiprocessor than in the uniprocessor because they overlap multiple load misses with similar multiprocessor (remote) latencies. The section of the code that exhibits overlap has a greater impact in the multiprocessor because of the longer remote latencies incurred in this section. d miss latency Simple ILP Simple ILP Radix209.2 overlapped stall | ||||||Utilization Number of L1 MSHRS Number of L1 MSHRS |||||||Utilization Number of L1 MSHRS Number of L1 MSHRS (a) Effect of ILP on average L1 miss latency (b) L1 MSHR occupancy due to loads (c) L1 MSHR occupancy due to loads and stores Fig. 6. Load miss overlap and contention in the ILP system. Fig. 7. ILP speedup for total execution time, CPU time, and memory stall time in the uniprocessor system. Consequently, the ILP multiprocessor generally sees lower parallel efficiency than the Simple multiprocessor. IV. Interaction of ILP Techniques with Software Prefetching The previous section shows that the ILP system sees a greater bottleneck from memory latency than the Simple system. Software-controlled non-binding prefetching has been shown to effectively hide memory latency in shared-memory multiprocessors with simple processors. This section evaluates how software prefetching interacts with ILP techniques in shared-memory multiprocessors. We followed the software prefetch algorithm developed by Mowry et al.[6] to insert prefetches in our applications by hand, with one exception. The algorithm in [6] assumes that locality is not maintained across synchronization, and so does not schedule prefetches across synchronization ac- cesses. We removed this restriction when beneficial. For a consistent comparison, the experiments reported are with prefetches scheduled identically for both Simple and ILP; the prefetches are scheduled at least 200 dynamic instructions before their corresponding demand accesses. The impact of this scheduling decision is discussed below, including the impact of varying this prefetch distance. A. Overall Results Figure graphically presents the key results from our experiments (FFT and FFTopt have similar performance, so only FFTopt appears in the figure). The figure shows the execution time (and its components) for each application on Simple and ILP, both without and with software prefetching ( +PF indicates the addition of software prefetching). Execution times are normalized to the time for the application on Simple without prefetching. Figure 9 summarizes some key data. Software prefetching achieves significant reductions in execution time on ILP (13% to 43%) for three cases (LU, Mp3d, and Water). These reductions are similar to or greater than those in Simple for these applications. How- ever, software prefetching is less effective at reducing memory stalls on ILP than on Simple (average reduction of 32% in ILP, ranging from 7% to 72%, vs. average 59% and range of 21% to 88% in Simple). The net effect is that even after prefetching is applied to ILP, the average memory stall time is 39% on ILP with a range of 11% to 65% (vs. average of 16% and range of 1% to 29% for Simple). For most ap- plications, the ILP system remains largely memory-bound even with software prefetching. B. Factors Contributing to the Effectiveness of Software Prefetching We next identify three factors that make software prefetching less successful in reducing memory stall time in ILP than in Simple, two factors that allow ILP additional benefits in memory stall reduction not available in Simple, and one factor that can either help or hurt ILP. We focus on issues that are specific to ILP systems; previous work has discussed non-ILP specific issues [6]. Figure 10 summarizes the effects that were exhibited by the applications we studied. Of the negative effects, the first two are the most important for our applications. Increased late prefetches. The last column of Figure 9 shows that the number of prefetches that are too late to completely hide the miss latency increases in all our applications when moving from Simple to ILP. One reason for this increase is that multiple-issue and out-of- order scheduling speed up computation in ILP, decreasing the computation time with which each prefetch is over- lapped. Simple also stalls on any load misses that are not prefetched or that incur a late prefetch, thereby allowing other outstanding prefetched data to arrive at the cache. ILP does not provide similar leeway. Increased resource contention. As shown in Section III, ILP processors stress system resources more than Simple. Prefetches further increase demand for resources, resulting in more contention and greater memory latencies. The resources most stressed in our configuration were cache ports, MSHRs, ALUs, and address generation units. Negative interaction with clustered misses. Optimizations to cluster load misses for the ILP system, as in LUopt, can potentially reduce the effectiveness of software prefetching. For example, the addition of prefetching re- dexecution time LU Memory CPU | ||||||Normalized execution time LUopt Memory CPU ||||||Normalized execution time FFTopt Memory CPU ||||||Normalized execution time Mp3d Memory CPU | ||||||Normalized execution time Radix Memory CPU ||||||Normalized execution time Water Memory CPU Fig. 8. Interaction between software prefetching and ILP. duces the execution time of LU by 13% on the ILP system; in contrast, LUopt improves by only 3%. (On the Simple system, both LU and LUopt improve by about 10% with prefetching.) LUopt with prefetching is slightly better than LU with prefetching on ILP (by 3%). The clustering optimization used in LUopt reduces the computation between successive misses, contributing to a high number of late prefetches and increased contention with prefetching. Overlapped accesses. In ILP, accesses that are difficult to prefetch may be overlapped because of non-blocking loads and out-of-order scheduling. Prefetched lines in LU and LUopt often suffer from L1 cache conflicts, resulting in these lines being replaced to the L2 cache before being used by the demand accesses. This L2 cache latency results in stall time in Simple, but can be overlapped by the processor in ILP. Since prefetching in ILP only needs to target those accesses that are not already overlapped by ILP, it can appear more effective in ILP than in Simple. Fewer early prefetches. Early prefetches are those where the prefetched lines are either invalidated or replaced before their corresponding demand accesses. Early prefetches can hinder demand accesses by invalidating or replacing needed data from the same or other caches without providing any benefits in latency reduction. In many of our applications, the number of early prefetches drops in ILP, improving the effectiveness of prefetching for these applications. This reduction occurs because the ILP system allows less time between a prefetch and its subsequent demand access, decreasing the likelihood of an intervening invalidation or replacement. Speculative prefetches. In ILP, prefetch instructions can be speculatively issued past a mispredicted branch. Speculative prefetches can potentially hurt performance by bringing unnecessary lines into the cache, or by bringing needed lines into the cache too early. Speculative prefetches can also help performance by initiating a prefetch for a needed line early enough to hide its latency. In our appli- cations, most prefetches issued past mispredicted branches were to lines also accessed on the correct path. App. % reduction in execution time reduction in memory stall time maining memory stall time prefetches that are late ple ple ple ple Mp3d 43 43 78 59 29 62 1 12 Water Average 14 14 59 Fig. 9. Detailed data on effectiveness of software prefetching. For the average, from LU and LUopt, only LUopt is considered since it provides better performance than LU with prefetching and ILP. Factor LU LU FFT Mp3d Water Radix opt opt Late prefetches Resource contention Clustered load Overlapped Early prefetches Speculative prefetches Fig. 10. Factors affecting the performance of prefetching for ILP. C. Impact of Software Prefetching on Execution Time Despite its reduced effectiveness in addressing memory stall time, software prefetching achieves significant execution time reductions with ILP in three cases (LU, Mp3d, and Water) for two main reasons. First, memory stall time contributes a larger portion of total execution time in ILP. Thus, even a reduction of a small fraction of memory stall time can imply a reduction in overall execution time similar to or greater than that seen in Simple. Second, ILP systems see less instruction overhead from prefetching compared to Simple systems, because ILP techniques allow the overlap of these instructions with other computation. D. Alleviating Late Prefetches and Contention Our results show that late prefetches and resource contention are the two key limitations to the effectiveness of prefetching on ILP. We tried several straightforward modifications to the prefetching algorithm and the system to address these limitations [12]. Specifically, we doubled and quadrupled the prefetch distance (i.e., the distance between a prefetch and the corresponding demand access), and increased the number of MSHRs. However, these modifications traded off benefits among late prefetches, early prefetches, and contention, without improving the combination of these factors enough to improve overall per- formance. We also tried varying the prefetch distance for each access according to the expected latency of that access (versus a common distance for all accesses), and prefetching only to the L2 cache. These modifications achieved their purpose, but did not provide a significant performance benefit for our applications [12]. V. Discussion Our results show that shared-memory systems are limited in their effectiveness in exploiting ILP processors due to limited benefits of ILP techniques for the memory sys- tem. The analysis of Section III implies that the key reasons for the limited benefits are the lack of opportunity for overlapping load misses and/or increased contention in the system. Compiler optimizations akin to the loop interchanges used to generate LUopt and FFTopt may be able to expose more potential for load miss overlap in an applica- tion. The simple loop interchange used in LUopt provides a 13% reduction in execution time compared to LU on an ILP multiprocessor. Hardware enhancements can also increase load miss overlap; e.g., through a larger instruction window. Targeting contention requires increased hardware resources, or other latency reduction techniques. The results of Section IV show that while software prefetching improves memory system performance with ILP processors, it does not change the memory-bound nature of these systems for most of the applications because the latencies are too long to hide with prefetching and/or because of increased contention. Our results motivate prefetching algorithms that are sensitive to increases in resource usage. They also motivate latency-reducing (rather than tolerating) techniques such as producer- initiated communication, which can improve the effectiveness of prefetching [1]. VI. Conclusions This paper evaluates the impact of ILP techniques supported by state-of-the-art processors on the performance of shared-memory multiprocessors. All our applications see performance improvements from current ILP techniques. However, while ILP techniques effectively address the CPU component of execution time, they are less successful in improving data memory stall time. These applications do not see the full benefit of the latency-tolerating features of ILP processors because of insufficient opportunities to overlap multiple load misses and increased contention for system resources from more frequent memory accesses. Thus, ILP-based multiprocessors see a larger bottleneck from memory system performance and generally poorer parallel efficiencies than previous-generation multiprocessors. Software-controlled non-binding prefetching is a latency hiding technique widely recommended for previous-generation shared-memory multiprocessors. We find that while software prefetching results in substantial reductions in execution time for some cases on the ILP system, increased late prefetches and increased contention for resources cause software prefetching to be less effective in reducing memory stall time in ILP-based systems. Even after the addition of software prefetching, most of our applications remain largely memory bound. Thus, despite the latency-tolerating techniques integrated within ILP processors, multiprocessors built from ILP processors have a greater need for additional techniques to hide or reduce memory latency than previous-generation multiprocessors. One ILP-specific technique discussed in this paper is the software clustering of load misses. Additionally, latency-reducing techniques such as producer-initiated communication that can improve the effectiveness of prefetching appear promising. --R An Evaluation of Fine-Grain Producer- Initiated Communication in Cache-Coherent Multiprocessors Adaptive and Integrated Data Cache Prefetching for Shared-Memory Multiprocessors The SGI Origin Tolerating Latency through Software-controlled Data Prefetching Evaluation of Design Alternatives for a Multi-processor Microprocessor The Case for a Single-Chip Multiprocessor An Evaluation of Memory Consistency Models for Shared-Memory Systems with ILP Processors RSIM Reference Manual The Impact of Instruction Level Parallelism on Multiprocessor Performance and Simulation Methodology. The Interaction of Software Prefetching with ILP Processors in Shared-Memory Systems SPLASH: Stanford Parallel Applications for Shared-Memory The SPLASH-2 Programs: Characterization and Methodological Considerations --TR --CTR Vijay S. Pai , Sarita Adve, Code transformations to improve memory parallelism, Proceedings of the 32nd annual ACM/IEEE international symposium on Microarchitecture, p.147-155, November 16-18, 1999, Haifa, Israel Manuel E. Acacio , Jos Gonzlez , Jos M. Garca , Jos Duato, Owner prediction for accelerating cache-to-cache transfer misses in a cc-NUMA architecture, Proceedings of the 2002 ACM/IEEE conference on Supercomputing, p.1-12, November 16, 2002, Baltimore, Maryland Christopher J. Hughes , Praful Kaul , Sarita V. Adve , Rohit Jain , Chanik Park , Jayanth Srinivasan, Variability in the execution of multimedia applications and implications for architecture, ACM SIGARCH Computer Architecture News, v.29 n.2, p.254-265, May 2001 Christopher J. Hughes , Sarita V. Adve, Memory-side prefetching for linked data structures for processor-in-memory systems, Journal of Parallel and Distributed Computing, v.65 n.4, p.448-463, April 2005 Xian-He Sun , Surendra Byna , Yong Chen, Server-based data push architecture for multi-processor environments, Journal of Computer Science and Technology, v.22 n.5, p.641-652, September 2007

instruction-level parallelism;performance evaluation;shared-memory multiprocessors;software prefetching

297811

Bayesian Function Learning Using MCMC Methods.

AbstractThe paper deals with the problem of reconstructing a continuous one-dimensional function from discrete noisy samples. The measurements may also be indirect in the sense that the samples may be the output of a linear operator applied to the function (linear inverse problem, deconvolution). In some cases, the linear operator could even contain unknown parameters that are estimated from a second experiment (joint identification-deconvolution problem). Bayesian estimation provides a unified treatment of this class of problems, but the practical calculation of posterior densities leads to analytically intractable integrals. In the paper it is shown that a rigourous Bayesian solution can be efficiently implemented by resorting to a MCMC (Markov chain Monte Carlo) simulation scheme. In particular, it is discussed how the structure of the problem can be exploited in order to improve computational and convergence performances. The effectiveness of the proposed scheme is demonstrated on two classical benchmark problems as well as on the analysis of IVGTT (IntraVenous Glucose Tolerance Test) data, a complex identification-deconvolution problem concerning the estimation of the insulin secretion rate following the administration of an intravenous glucose injection.

Introduction The problem of reconstructing (learning) an unknown function from a set of experimental data plays a fundamental role in engineering and science. In the present paper, the attention is restricted to scalar functions of a scalar variable the main ideas may apply to more general maps as well. In the most favourable case, it is possible to directly sample the function. Rather frequently, however, only indirect measurements are available which are obtained by sampling the output of a linear operator applied to the function. For instance, in deconvolution problems what is sampled is the convolution of the unknown function with a known kernel, see e.g. [1], [2], [3], [4], [5]. The approaches used to solve the function learning problem can be classified according to three major strategies. The parametric methods assume that the unknown function belongs to a set of functions which are parameterized by a finite-dimensional parameters vector. For instance, the function can be modelled as the output of a multi-layer perceptron which, for a given topology, is completely characterized by the values of its weights [6]. Another possibility is to use a polynomial spline with fixed knots. In both cases, function learning reduces to the problem of estimating the model parameters, a task that can be performed by solving a (possibly nonlinear) least squares problem. If it were true that the unknown function belongs to the given function space, statistical estimation theory could be invoked in order to find minimum variance estimators and compute confidence intervals [7]. Moreover, it would also be possible to compare parametric models of increasing complexity using statistical tests (F-test) or complexity criteria (Akaike's criterion). The second strategy, namely regularization [8], [9], [10], [11], avoids introducing heavy assumptions on the nature of F(\Delta) but rather classifies the potential solutions according to their "regularity" (typically by using an index of smoothness such as the integral of the squared k-th derivative of the function). The relative importance of the sum of squared residuals against the regularity index is controlled by the so-called regularization parameter. The key problem is finding an optimal criterion for the selection of the regularization parameter, although empirical criteria such as ordinary cross validation and generalized cross validation [12], [13] perform satisfactorily in many practical cases. Moreover, regularization does not provide confidence intervals so that it is not possible to assess the reliability of the reconstructed function. The present paper deals with the third strategy which is based on Bayesian estimation. The unknown function is seen as an element of a probability space whose probability distribution reflects the prior knowledge. For instance, the prior knowledge that F(\Delta) is smooth is translated in a probability distribution that assigns higher probabilities to functions whose derivatives have "small" absolute values. A practical way to do that is to describe F(\Delta) as a Gaussian stochastic process whose k-th derivative is a white noise process with intensity - 2 . Provided that both - 2 and the variance oe 2 of the (Gaussian) measurement noise are known, the Bayes formula can be used to work out the posterior distribution of F(\Delta) given the data [2], [14], [15], [16]. The posterior provides a complete description of our state of knowledge. In particular, the mean of the posterior can be used as a point estimate (Bayes estimate) whereas the variance helps assessing the accuracy. It is notable that, if the regularization parameter is taken equal to the ratio oe 2 =- 2 the regularized estimate coincides with the Bayes one. The main advantage of the Bayesian approach is the possibility to address the selection of the regularization parameter in a rigourous probabilistic framework. In fact, when - 2 is not known, it can be modelled as a random variable and two different approaches are possible. The simpler one is based on the following observation: if the prior distribution of - 2 is very flat, the maximum of its posterior given the data is close to the maximum likelihood estimate - 2 ML . Then, if the posterior of - 2 is very peaked around its max- imum, it is reasonable to estimate F(\Delta) as if - 2 ML were the true value of - 2 [14], [17], [4]. MAGNI ET AL.: BAYESIAN FUNCTION LEARNING USING MCMC METHODS 3 However, if the posterior of - 2 is not "very peaked" (which is likely to happen especially for medium and small data sets), neglecting the uncertainty on - 2 would lead to under-estimated confidence margins for F(\Delta). The truly Bayesian approach, conversely, calls for the computation of the posterior of F(\Delta) taking into account also the random nature of . Since the involved integrals are analytically intractable, one has to resort to Monte Carlo methods. A first purpose of the present paper is to show how a truly Bayesian solution of the function learning problem can be efficiently worked out. We discuss the various stages of the procedure starting from the discretization of F(\Delta) to arrive at the practical implementation of Markov chain Monte Carlo (MCMC) methods [18]. Our approach is similar to the one proposed in [15] where, however, the case of indirect measurements (deconvolution problem) is not treated. A further issue addressed in the paper is the joint identification-deconvolution problem, which arises when the convolution kernel is not a priori known but is to be identified by a separate experiment. The standard (suboptimal) approach is to identify the convolution kernel and then use it as if it were perfectly known in order to learn the unknown function F(\Delta). As shown in the paper, the use of MCMC methods allows to learn both functions jointly. The paper is organized as follows. Section II contains the statement of the problem. In Section III, after a concise review of MCMC methods, a numerical procedure for solving the Bayesian function learning problem is worked out. In section Section IV the proposed method is illustrated by means of simulated as well as real-world data coming from the analysis of metabolic systems. Some conclusions (Section V) end the paper. II. Problem statement In this paper we consider the problem of reconstructing a function ~ discrete and noisy samples y k such that ~ denotes the measurement error and ~ L k is a linear functional. In increasing order of generality, we have: ~ ~ ~ ~ denote the sampling instants and t is the initial time. The first definition of ~ corresponds to the function approximation problem based on samples of the function itself. Using the second and third definitions, the problem (1) becomes a deconvolution problem, or an integral equation of the first kind (also called Fredholm equation), respectively. As for the noise v k , letting v := [v 1 assumed that is a known matrix and oe 2 is a (possibly unknown) scalar. In the following, y := [y 1 ~ f ~ ~ non parametric estimator based on Tychonov regularization is given by ~ ~ f P is a suitable operator and k \Delta k is a norm in a suitable function space. A typical choice for ~ is: ~ Then, if in addition ~ L k is as in (2) and "L 2 norm" is used, ~ turns out to be a smoothing spline [13]. The basic idea behind (5) is to find a balance between data fit and smooth- MAGNI ET AL.: BAYESIAN FUNCTION LEARNING USING MCMC METHODS 5 ness of the solution, the relative weight being controlled by the so-called regularization parameter fl. Various criteria have been proposed to tune fl. Among them, we may mention ordinary cross-validation [1], generalized cross-validation [12], and the L-curve [19]. It is worth noting that without the aid of the smoothness penalty flk ~ would be impossible to learn the function ~ f (belonging to an infinite dimensional space) from the finite data set fy k g. Rather interestingly, (5) can also be interpreted as a Bayesian estimator. To this pur- pose, assume that ~ f is a stochastic process such that ~ f is a white noise with intensity ~ w(t)]=0, 8t and E[ ~ w(t) ~ assuming that both v k and ~ f are normally distributed and letting , the regularized estimate ~ coincides with the conditional expectation of ~ f given the observations y k , i.e. E[ ~ f fl . According to the Bayesian paradigm, the probabilistic assumptions on the unknown function ~ f should reflect our prior knowledge. For instance, if the operator ~ P is as in (6), then ~ f is the double integral of a white noise: as such it is a relatively smooth signal (the smaller ~ the smoother ~ f will be). The above considerations suggest that ~ f fl is the "optimal" estimator provided that . Unfortunately, ~ - 2 (and sometimes also oe 2 ) is very unlikely to be known a priori. In the literature [20], [14], [4] it has been proposed to regard ~ as an unknown parameter and compute its maximum likelihood estimate: ~ In this way ~ f fl with becomes a sub-optimal estimator. In this paper, conversely, we pursue a rigourous Bayesian approach. The unknown parameter ~ - 2 is modelled as a hyper-parameter having a suitable prior distribution, which is taken into account in the computation of the posterior density Z ~ f)p( ~ 6 IEEE TRANSACTIONS ON PATTERN ANALYSIS AND MACHINE INTELLIGENCE Since the analytic evaluation of the posterior is made intractable by the presence of the hyper-parameter, the calculations are carried out by means of Monte Carlo sampling algorithms. In this context it is convenient to discretize the original problem. Without loss of generality, it is assumed that t 0 =0. Then, given a sufficiently small discretization interval T , consider f . For simplicity, it is assumed that the sampling instants t k are multiples of T , i.e. t the operator ~ P is suitably discretized. For instance, with approximated by: Moreover, ~ f is approximated by Lf , where L is a suitable n \Theta N matrix: if ~ L k is as in (2) then L kj if ~ L k is as in (3) then L R T~ if ~ L k is as in (4) then L R T~ These approximations are obtained under the assumption that the unknown function ~ f is constant in between the sampling instants. After the discretization, (5) is approximated by: f denotes the usual norm in ! n ), whose closed form solution is: MAGNI ET AL.: BAYESIAN FUNCTION LEARNING USING MCMC METHODS 7 In the Bayesian setting it is assumed that the entries of the vector are taken from the realization of a (discrete-time) white noise with variance - 2 I, that is E[w]= 0, A first purpose of the present paper is to investigate the applicability of Markov chain Monte Carlo [18] sampling methods to estimate p(f j y) assuming that - 2 is a parameter with suitable prior distribution. Another issue concerns the Bayesian solution of the joint identification-deconvolution problem. In fact, in many deconvolution problems the impulse response ~ h(-) that enters the convolution integral (3), depends on one or more parameters -, that are estimated by a separate experiment: ~ putting together the two experiments, we obtain: where M(-) is a known function of -, z := [z is the data vector used to identify - and ffl := is the corresponding measurement error. The standard approach is to estimate - using (10) and then estimate f from using - as if it were the true value of -. On the other hand, a truly Bayesian approach describes - as a random variable. Then, p(f j y) should be evaluated by considering (9) and (10) simultaneously. As a particular case, it is possible to consider (9) alone with - modelled as a random variable to allow for its uncertainty. Again, the standard "suboptimal" approach is to compute f fl using the nominal value of - [4] and then assess the sensitivity of the estimate with respect to parameters uncertainty. III. Markov chain Monte Carlo Methods in Bayesian estimation problems Probabilistic inference involves the integration over possibly high-dimensional probability distributions. Since this operation is often analytically intractable, it is common to 8 IEEE TRANSACTIONS ON PATTERN ANALYSIS AND MACHINE INTELLIGENCE resort to Monte Carlo techniques, that requires sampling from the probability distribution to be integrated. Unfortunately, sometimes it is impossible to extract samples directly from that distribution. Markov chain Monte Carlo (MCMC) methods [18] provide a unified framework to solve this problem. MCMC methods are based on two steps: a Markov chain and a Monte Carlo integra- tion. By sampling from suitable probability distributions, it is generated a Markov chain that converges (in distribution) to the target distribution, i.e. the distribution to be in- tegrated. Then, the expectation value is calculated through Monte Carlo integration over the obtained samples. The MCMC methods differ from each other in the way the Markov chain is generated. However, all the different strategies proposed in the literature, are special cases of the Metropolis-Hastings [21], [22] framework. Also the well-known Gibbs sampler [23] fits in the Metropolis-Hastings scheme. In the following sub-sections, we will describe the application of the Metropolis-Hastings algorithm to Bayesian function learning as well as discuss about its possible variants. A. The method In order to describe the Metropolis-Hastings algorithm we will use the following notation the vector of the model parameters the vector of the data (i.e. the observations) ffl \Theta i - the i th samples drawn the target distribution (proportional to the posterior distribution) where p(') is the prior distribution of the model parameters and p y the likelihood. MAGNI ET AL.: BAYESIAN FUNCTION LEARNING USING MCMC METHODS 9 The Markov chain derived by the Metropolis-Hastings method is obtained through the following steps: 1. at each time t, a candidate sample \Theta is drawn from a proposal distribution 2. the candidate point \Theta is accepted with probability: 3. if the candidate point \Theta is accepted, the next sample of the Markov chain is \Theta else the chain does not move and \Theta It is important to remark that the stationary distribution of the chain (i.e. the distribution to which the chain converges) is independent of the proposal distribution [18], and coincides with the target distribution p ';y ('). Although any proposal distribution, on long-run, will deliver samples from the target distribution, the rate of convergence to the stationary distribution of the generated Markov chain crucially depends on the relationships between the proposal and the target distributions; moreover, the number of samples necessary to perform the Monte Carlo steps depends on the speed with which the algorithm "mixes" (i.e. spans the support of the target distribution). When the vector of the model parameters is large, it is often convenient to divide ' into K components and update the samples \Theta of these components one by one [24]. This scheme is called Single-Component Metropolis-Hastings Let \Theta (i) t the i th component of \Theta at time t and \Theta (\Gammai) \Theta (K) t g, the Metropolis-Hastings scheme turns to: 1. at each time t, the next sample \Theta (i) t+1 is derived by sampling a candidate point \Theta (i) from a proposal distribution q (i) 2. the candidate point \Theta (i) is accepted with probability: ff(\Theta (i) 3. if the candidate point \Theta (i) is accepted, the next sample of the Markov chain is \Theta (i) \Theta (i) , else the chain does not move and \Theta (i) t . The Gibbs sampler (GS) is just a special case of the single-component Metropolis- Hastings. The GS scheme exploits the full conditional (the product of the prior distribution and the likelihood) as the proposal distribution. In this case, it is easy to verify that the candidate point is always accepted, so that the Markov chain moves at every step. When the full conditionals are standard distributions (easy to sample from), the GS represents a suitable choice. On the contrary, when it is not possible to draw samples directly from the full-conditional distributions, it is convenient to resort to mixed schemes Metropolis-Hastings). In this setting, a portion of the model parameters is estimated using the Gibbs Sampler, while the other ones are treated using "ad-hoc" proposal distributions. These algorithms have been extensively used in the field of probabilistic graphical modelling [25]. Using this kind of models a suitable partition (blocking) of the vector of model parameters is naturally obtained and it is also easy to derive the full conditional distribu- tion. The convergence rate and the strategies for choosing the proposal distribution are described in [26], [27]. B. MCMC in Function Reconstruction In this sub-section we will describe how the problems defined in Section II can be tackled using MCMC methods. In order to explain the probabilistic models used in the different sampling schemes, we will resort to a Bayesian Network (BN) representation. BNs are Directed Acyclic Graphs (DAGs) in which nodes represent variables, while arcs express direct dependencies between variables. These models are quantified by specifying the conditional probability distribution of each node given its parents. They will help us in MAGNI ET AL.: BAYESIAN FUNCTION LEARNING USING MCMC METHODS 11 expressing the conditional independence assumptions underlying the different function reconstruction problems. For further details see [28], [29]. B.1 Function approximation based on direct sampling Consider the function approximation problem based on samples of the function itself (smoothing problem). Its formal definition is given in equations (1) and (2). Our goal is to provide a Bayesian estimate of the vector f (i.e. the discretized unknown function). As shown in Section II, the discretized form of the problem may be written as: Referring to Section II, f :=P \Gamma1 w with E[w]=0, E[ww T To apply the MCMC strategy described in Section III-A, we must assign a suitable probabilistic model to the parameter set g. By exploiting a set of standard choices in Bayesian estimation problems, it is assumed that w is normally distributed given - 2 and that the precision parameter 1 has a Gamma distribution. More formally, and . Moreover, we suppose that the noise v has a Normal distribution, covariance matrix oe 2 \Psi. This implies that the data model can be written as: p(y In this setting the target distribution becomes: The model is described by the simple BN of Fig. 1 It is easy to see that, in order to apply MCMC integration, it is useful to adopt the partition =wg. In this context, it is convenient to adopt the Gibbs Sampler, since the full conditional distributions assume the following standard form: y I Y Fig. 1. Function approximation: probabilistic model of the smoothing problem, deconvolution problem and Fredholm equation problem. where From the point estimate - derived by the MCMC algorithm, it is trivial to reconstruct the unknown function f as - Moreover, having samples from the joint posterior distribution, it is possible to derive any statistics of interest, including confidence intervals (or more appropriately: Bayes intervals). B.2 Function approximation in inverse problems In deconvolution problems the unknown function has to be reconstructed on the basis of indirect measurements: a convolution integral expresses the relationships between the samples and the unknown function, see (1) and (3). The structure of the problem is analogous when considering integral equations of the first kind (Fredholm equations), see (1) and (4) Again, our goal is to provide a Bayesian estimate of the (discretized) unknown f func- MAGNI ET AL.: BAYESIAN FUNCTION LEARNING USING MCMC METHODS 13 tion. The discretized form of the problem still has the form (11). Since the functions ~ h(-) and ~ h(t; -) that enter in the definitions (3) and (4) are assumed to be known, the matrix L is completely specified. Also in this case, f := P \Gamma1 w with so that the parameter set to be estimated is the same as in the smoothing problem: g. Thus, the probabilistic model for the inverse problems (1),(3) and (1),(4) is again described by the BN of Fig. 1. In fact, in our setting, the smoothing problem, the deconvolution problem and the Fredholm equation problem differ only in the computation of the matrix L. B.3 Deconvolution problems with uncertain impulse response An interesting extension of the problem described in the previous section is to relax the assumption of complete knowledge of the impulse response ~ h(-) appearing in equation (3). As anticipated in Section II, we suppose that ~ is a function of a set of unknown parameters -, which have to be estimated from experimental data. In this way, the problem becomes the simultaneous estimation of the unknown function f and the parameter set -, given the model described by (9) and (10). Once again, f := P \Gamma1 w with I. In this case, however, the probabilistic model to be specified becomes more complex. The parameter set to be estimated is now -g. The corresponding probabilistic model is assumed as: - 0 and \Sigma - . Moreover, we suppose that the noise signals v and ffl are independent and normally distributed with known covariance matrices, namely v - N(0; oe 2 \Psi) and ffl - Then, the data model can be written as: p(y j w; In this setting the target distribution becomes: The resulting model is described by the BN of Fig. 2. 14 IEEE TRANSACTIONS ON PATTERN ANALYSIS AND MACHINE INTELLIGENCE y I x x x z ~ N(M( 1/ l YFig. 2. Function approximation in deconvolution problems in presence of uncertainty on the impulse response: probabilistic model. In order to devise an MCMC scheme for estimating ', it is convenient to partition ' as =-g. The full conditional distributions are as follows: where and The full conditional (16) is clearly a non-standard distribution, so that it cannot be sampled directly. In order to use the GS strategy it is necessary to apply sampling algorithms for general distributions, like rejection sampling or adaptive rejection sampling MAGNI ET AL.: BAYESIAN FUNCTION LEARNING USING MCMC METHODS 15 [30]. Unfortunately, such algorithms may impair the overall efficiency of the stochastic simulation machinery. A valuable alternative is to resort to a mixed MCMC scheme in which different proposal distributions are used to extract samples from the different partitions of '. ffl The proposal for ' (1) ; ' (2) are the full conditional distributions, as in the GS. In this way the candidate point for these partitions is always accepted. ffl The proposal distribution for ' (3) can be chosen as the prior distribution In this case, the proposal distribution is independent from the past sample drawn by the Markov chain, so that This scheme is also known as independence sampler [31]. The acceptance probability for the candidate sample \Theta (3) simplifies as: ff(\Theta (3) so that the acceptance rate depends only on the ratio of the likelihood in the candidate point to the likelihood for the current one. On the basis of the specific problem, other proposal distributions can be chosen in order to obtain the best performance in the computational speed. For example, when possible, a good choice can be the use of a standard distribution that is a good approximation of the full conditional distribution. This is a way to preserve the advantages of the single-component Metropolis-Hastings, avoiding the additional computational burden that would be entailed by the GS in presence of non-standard full conditional distributions. IV. Bayesian function learning at work In this section we will show how the above presented methodology is able to cope with three different benchmark problems taken from the literature. A. Function approximation based on direct sampling To test the performance of the MCMC function approximator in the smoothing problem, we consider an example proposed by Wahba [13]. The function to be approximated is: ~ and the noisy samples y k are: Since we are interested in reconstructing the function only in correspondence of the measurement times, in equation (11) L is the identity matrix. Following Section III-B.1 we take the following prior where I is the identity matrix. The choice of the prior distribution parameters for the- 2 reflects the absence of reliable prior information on the regularization parameter. In fact with this choice - with probability of 0.9. As in [13], we assume that the second derivative of the function is regular; taking into account the discretization, the operator P is chosen as: MAGNI ET AL.: BAYESIAN FUNCTION LEARNING USING MCMC METHODS 17 -0.50.5t (arbitrary units) f (arbitrary units) Fig. 3. Function approximation using the MCMC-smoother.: noisy data (stars), true function (dash-dot line) and reconstructed function (continuous line). The starting point of the Markov chain was extracted from the prior distribution of the parameters fw; - 2 g. After a 1000 samples run of the MCMC scheme, convergence of the estimates was verified by the method described in [26]. In particular, after choosing the quantiles 0:975g to be estimated with precision respectively with probability burn-in (N) of samples and a number of required samples (M) of 870 were calculated. The results are shown in Fig. 3. Although the samples are rather noisy, the smoothed signal is close to the true function. The RMSE (root mean square error) obtained is 0.073. Moreover, the performances of the MCMC smoother are similar to the ones obtained in [13], where a cubic smoothing spline was used and the regularization parameter was tuned through ordinary cross-validation (OCV). Hence, the MCMC smoother is as good as OCV-tuned smoothing splines in avoiding under- and over-smoothing problems. The main advantage of MCMC smoother is that it provides also the a-posteriori sampling distribution of the regularization parameter and the confidence intervals for the reconstructed function in a rigourous Bayesian setting. B. Function approximation in deconvolution problems In order to test the performances of the MCMC deconvolution scheme, we consider a well-known benchmark problem [32], [2], [4]. The input signal given by: ~ has been convoluted with the impulse response: Then, by adding measurement errors v k simulated as a zero-mean white Gaussian noise sequence with variance equal to 9, 52 noisy samples are generated at time t Fig. 4(a) shows the true function ~ while Fig. 4(b) depicts the convoluted function together with the 52 noisy samples. Our goal is to reconstruct the unknown function with a sufficiently fine resolution. This means that we are interested in the function estimates not only in correspondence of the measurement times, but also in other "out of samples" time points. In particular, we consider a 208 points on an evenly-spaced time grid in the interval [0 1035]. Then, the entries of L are R T~ with We take the following prior distributions, similar to the ones used in the previous section:- 2 - \Gamma(0:25; 5e \Gamma 7) MAGNI ET AL.: BAYESIAN FUNCTION LEARNING USING MCMC METHODS 19 (a) (arbitrary units) f (arbitrary units) (arbitrary units) y (arbitrary units) Fig. 4. Simulated deconvolution problem. Panel (a): the function to be reconstructed. Panel (b): the convoluted function (dash-dot line) and the noisy samples (stars). Moreover the operator P is selected as: This choice corresponds to a penalty on the squared norm of the first derivative, (which is approximated by the first difference of the discretized signal). The bottleneck of the MCMC scheme is the computation of B \Gamma1 in equation (15)(this matrix inverse has to be performed at each step of the MCMC scheme). However, by a proper change of coordinates, it is possible reduce the size of the matrix has to be inverted from N \Theta N to n \Theta n (i.e. from 208 \Theta 208 to 52 \Theta 52 in our problem). This goal can be achieved through the following steps: 1. Let hence an n \Theta N matrix), and compute the SVD (singular value n) and V (N \Theta N) are orthogonal matrices (UU and D is an n \Theta N diagonal matrix (D 2. In view of (11) It is easy to verify that v is distributed as N(0; I n ), and w as N(0; - 2 I N ). 3. We apply the same MCMC scheme described in the previous section to the reformulated problem (17). In the new coordinates B is a block diagonal matrix, in which the first one is an N \Theta N block and the other ones are 1 \Theta 1 blocks. 4. The final estimate is obtained by re-transforming the variables in the original co- ordinates; in particular we need to compute: The starting point of the Markov chain was extracted from the prior distribution of the parameters fw; - 2 g. After 5000 steps of the MCMC scheme, the convergence of the estimates was verified by using the method described in [26]. In particular after choosing the quantiles 0:975g to be estimated with precision respectively with probability burn-in (N) of 132 samples and a number of required samples (M) of 3756 were calculated. The results are shown in Fig. 5. The performance of our approach is comparable with the one proposed in [4], where the regularization parameter is estimated according to a maximum likelihood criterion (see Fig. 6, Fig. 7). MAGNI ET AL.: BAYESIAN FUNCTION LEARNING USING MCMC METHODS 21 (arbitrary units) f (arbitrary units) (arbitrary units) y (arbitrary units) Fig. 5. Simulated deconvolution problem. Panel (a): the true function (dash-dot line) and the re-constructed one (continuous line), with its 95% confidence interval (dashed line), obtained with the MCMC scheme. Panel (b): the true noiseless output (dash-dot line), the estimated output (continu- ous line) and the noisy samples (*). The RMSE obtained with our approach is 0.065 while the one obtained by the method of [4] is 0.059. Again, the advantage of the MCMC scheme is its ability to provide the a-posteriori sampling distribution of the regularization parameter and the confidence intervals for the reconstructed function in a rigourous Bayesian setting. C. Deconvolution with uncertain impulse response In this subsection, the MCMC scheme is applied to a real-world problem taken from [33]; in particular, we demonstrate that deconvolution and impulse response identification can be addressed jointly, as described in Section III-B.3. The goal is to quantify the Insulin Secretion Rate (ISR) in humans after a glucose stim- ulus; the experimental setting is related to the so-called IntraVenous Glucose Tolerance 22 IEEE TRANSACTIONS ON PATTERN ANALYSIS AND MACHINE INTELLIGENCE (arbitrary units) f (arbitrary units) Fig. 6. Simulated deconvolution problem. Comparison of the MCMC scheme with the maximum likelihood one (see text). True function (dashed line), MCMC estimate (continuous line), maximum likelihood estimate (thick line). Test (IVGTT), where an impulse dose of glucose is administered in order to assess the subject capability of bringing Blood Glucose Levels within normal ranges through the endogenous release of Insulin, the main glucoregolatory hormone. The ISR cannot be directly measured since insulin is secreted by the pancreas into the portal vein which is not accessible in vivo. It is possible measure only the effect of the secretion in the circulation (the plasma concentration of insulin). But, because of the large liver extraction, the plasma insulin concentration reflects only the post-hepatic delivery rate into the circulation. This problem can be circumvented by measuring C-peptide (CP) concentration in plasma. The CP is co-secreted with insulin on an equimolar basis, but is not extracted by liver, so that it directly reflects the pancreatic ISR. Thus, the problem turns into the estimation of the ISR on the basis of the (noisy) measurements of CP in plasma. Since the CP kinetics can be described by a linear model, we obtain the following MAGNI ET AL.: BAYESIAN FUNCTION LEARNING USING MCMC METHODS 23 samples lambda^2 samples of lambda^2 (b) Fig. 7. Simulated deconvolution problem. Panel (a): Markov chain of the - 2 parameter. Panel (b): frequency histogram proportional to the posterior distribution of the parameter - 2 estimated. The derived with the maximum likelihood approach (see text) was 0.00665. ~ where y k are the CP plasma measurements (pmol/ml), ~ f(\Delta) is the ISR (pmol/min), ~ h(\Delta) is the CP impulse response (ml \Gamma1 ) and v k is the measurement error (pmol/ml). The CP impulse response is [33]: ~ h(-) =X where the parameters A i s and ff i s have to be estimated through an ad-hoc experiment; in particular an intravenous bolus of biosynthetic CP is delivered to the patient and a number of plasma measurements of CP are collected (a somatostatin infusion is administered in order to avoid the endogenous pancreatic secretion). The impulse response parameters depend on the single patient but are considered constant over time for a specific patient. We can apply the MCMC scheme of Section III-B.3 in order to jointly perform the CP impulse response identification and the ISR reconstruction. Following Section II the problem can be written as: is the matrix obtained from the discretization of the deconvolution integral, z are the noisy measurements of the CP plasma concentration during the impulse response identification experiment, v and ffl are the measurements errors and are the sampling instants of the identification experiment. L(-) has elements: In this case, some prior knowledge on the signal is available [33]: the ISR is known to exhibit a spike just after the external glucose stimulus and a more regular profile thereafter. This knowledge can be modelled by considering two different regularization parameters for the two phases. For computational reasons, a nonuniform discretization for f has been adopted. Ac- cordingly, the regularization operator P has been chosen so as to satisfy: -T Then: MAGNI ET AL.: BAYESIAN FUNCTION LEARNING USING MCMC METHODS 25 y ~ N(0, L) x x ~ N( ,S x x z ~ N(M( 1/ l e s Fig. 8. Probabilistic model for the reconstruction of Insulin Secretion Rate Moreover, the variances oe 2 v and oe 2 ffl of the measurements errors v and ffl are only imprecisely known and must be estimated as well. The complete model is shown in Fig. 8. The derived sampling distributions for the MCMC scheme described in Section III-B.3 are: 26 IEEE TRANSACTIONS ON PATTERN ANALYSIS AND MACHINE INTELLIGENCE where In the above formulas N 1 is the number of point used in the regularization of the first region (the spike) while N 2 is the number of point used for the second region the vector w is such that the number of measurements in the identification experiment. is a diagonal matrix with the first N 1 elements equal to - 2 1 and with the other N 2 elements equal to - 2 2 . Finally, \Psi and \Psi ffl are diagonal matrices with elements y 2 and z 2 respectively; in this way, the parameters oe v and oe ffl represent the CV (coefficient of variation) of the measurements errors in the two experiments. To completely specify the MCMC scheme the following prior distributions must be 3 0:25 0:125 0:05] T and \Sigma - is a diagonal matrix with elements such that \Sigma - (i; MAGNI ET AL.: BAYESIAN FUNCTION LEARNING USING MCMC METHODS 27 The value of - 0 was derived on the basis of the prior knowledge on the dynamics of the impulse response, while the value of the hyper-parameters for oe v and oe ffl are assessed by knowing that the measurements error has a CV that ranges from 4% to 6%. The priors for the -s reflect the knowledge on the signal shape in the two response phases. We perform our test on the data set described in [34], [33]. The data set used for the impulse response identification are collected after a CP bolus of 49650 pmol, while the data set used for the ISR reconstruction are taken after an intravenous glucose bolus of 0:5 g=kg. The basal value of the ISR is estimated on the basis of the five CP measurements taken before the glucose bolus. Our goal is to reconstruct the ISR in correspondence of the sampling instants in which CP measurements are taken (in this case N=n). As in [33], we take N corresponding to a first phase ranging 16. The data of the two experiments are shown in Fig. 9. After a 2500 samples run of the MCMC scheme (convergence was verified by using the method described in [26] as the same assumption on q, r, s previous reported), the results shown in Fig. 10 were obtained. Fig. 10(a) shows the estimate of the CP plasma levels in the IVGTT experiment: the estimated curve is slightly smoother than the measurements. Fig. 10(b) depicts the ISR curve as estimated after deconvolution by the MCMC scheme: the reconstructed ISR reproduces the expected physiological shape, characterized by two regions with different regularities. The results obtained are comparable with those obtained in [33], where deconvolution and impulse response identification are treated separately. Fig. 10(c) shows the estimated CP impulse response identification. It is easy to notice the good quality of the fit. In Fig. 11 the frequency histograms of the samples generated by the MCMC estimator for the six impulse response parameters are reported. The proposed MCMC scheme is able to jointly perform the identification of the impulse response and the deconvolution of the ISR. In the classical approach [33], the two experi- 28 IEEE TRANSACTIONS ON PATTERN ANALYSIS AND MACHINE INTELLIGENCE (a) time (min) in plasma (pmol/ml) time (min) in plasma (pmol/ml) Fig. 9. The data set for the ISR (Insulin Secretion Rate) reconstruction. Panel (a): impulse response identification experiment, consisting of 32 noisy samples of the CP (C-Peptide concentration) in plasma, collected after a CP intravenous bolus at time 0 min. Panel (b): Intravenous Glucose Tolerance Test, consisting of 27 noisy samples of CP in plasma, collected around an intravenous glucose bolus at time 0 min. ments are treated in a separate fashion: in the first step, the impulse response is identified, using the measurements of the "identification set", and only in the second step, the data of the "deconvolution set" is used to reconstruct the unknown function. The uncertainty on the impulse response is possibly taken into account only after the deconvolution step. On the contrary, our scheme combines together the information coming from the two ex- periments, and uses it in order to provide "optimal" point estimates as well as posterior moments and confidence intervals. V. Conclusions MCMC methods constitute a set of emerging techniques, that have been successfully applied in a variety of contexts, from statistical physics to image analysis [23], and from MAGNI ET AL.: BAYESIAN FUNCTION LEARNING USING MCMC METHODS 29 time (min) in plasma (pmol/ml) (b) time (min) ISR (c) time (min) in plasma (pmol/ml) Fig. 10. Solution of the joint deconvolution and impulse response identification problem. Panel (a): estimated level of CP in plasma during the IVGTT (continuous line) and the collected samples (stars). Panel (b): the ISR reconstructed (with 95% confidence interval) by the MCMC-estimator. Panel (c): the estimated impulse response ~ h(\Delta) (continuous line) and the samples (stars) collected during the identification experiment. medical monitoring [25] to genetics [35] and archaeology[36]. The powerfulness of MCMC methodologies lies in their inherent generality that enables the analyst to deal with the full complexity of real world problems. In this paper, we have exploited such generality to propose a unified Bayesian framework for the reconstruction of functions from direct or indirect measurements. In particular, by using the same conceptual scheme we easily coped with problems that had been previously solved with ad-hoc methods. The obtained results are, in all cases, at least as samples of (a) samples of alpha2 (b) samples of alpha3 (c) value (1/ml) samples of (d) samples of samples of A3 Fig. 11. Frequency histograms proportional to the posterior distribution of estimated impulse response parameters. good as the previously proposed solutions. In addition, since our approach is able to soundly estimate the posterior probability distribution of the reconstructed function, the information provided at the end of the estimation procedure is richer than in all other methods: first and second moments, confidence intervals and posterior distributions are obtained as a by-product. Finally, our framework has been exploited to implement a new strategy for the joint estimation of a deconvoluted signal and its impulse response. The previous approaches were based on a two-step procedure, which is not able to optimally combine all the information available in the data. The main limits of MCMC approach is the time required to converge to the posterior distribution and the difficulty to choose the best sampling scheme. These limitations force to use MCMC methods only for off-line reconstructon. In summary, MCMC methods have been shown to play a crucial role in the off-line MAGNI ET AL.: BAYESIAN FUNCTION LEARNING USING MCMC METHODS 31 function learning problem, since they provide a flexible and relatively simple strategy, able to provide optimal results in a Bayesian sense. Acknowledgements The authors would like to thank Antonietta Mira for her methodological support in designing the MCMC scheme and Claudio Cobelli and Giovanni Sparacino for having provided the experimental data for the Insulin Secretion Rate reconstruction problem. They thank also the anonymous reviewers for their useful suggestions. --R "Practical approximate solutions to linear operator equations when the data are noisy," "The deconvolution problem: Fast algorithms including the preconditioned conjugate-gradient to compute a map estimator," "Linear inverse problems and ill-posed problems," "Nonparametric input estimation in physiological systems: Problems, methods, case studies," "Blind deconvolution via sequential imputations," Bayesian Learning for Neural Networks Parameter Estimation in Engineering and Science "A technique for the numerical solution of certain integral equations of the first kind," "On the numerical solution of Fredholm integral equations of the first kind by the inversion of the linear system produced by quadrature," Solutions of Ill-Posed Problems "Networks for approximation and learning," "Smoothing noisy data with spline functions: Estimating the correct degree of smoothing by the method of generalized cross-validation," Spline Models for Observational Data "Bayesian interpolation," "Gaussian processes for regression," "Automatic bayesian curve fitting," "Nonparametric spline regression with prior information," Markov Chain Monte Carlo in Practice "Numerical tools for analysis and solution of Fredholm integral equation of the first kind," "A time series approach to numerical differenti- ation," "Equations of state calculations by fast computing machine," "Monte Carlo sampling methods using Markov Chain and their applications," "Stochastic relaxation, Gibbs distributions, and the bayesian restoration of images," "Likelihood analysis of non-gaussian measurment time series," "A unified approach for modeling longitudinal and failure time data, with application in medical monitoring," "Implementing MCMC," "Inference and monitoring convergence," Probabilistic Reasoning in Intelligent Systems "Dynamic probabilistic networks for modelling and identifying dynamic systems: a MCMC approach," "Adaptive rejection sampling for Gibbs sampling," "Markov Chains for posterior distributions (with discussion)," "The inverse problem of radiography," "A stochastic deconvolution method to reconstruct insulin secretion rate after a glucose stimulus," "Peripheral insulin parallels changes in insulin secretion more closely than C-peptide after bolus intravenous glucose administration," "Censored survival models for genetic epidemi- ology: a Gibbs sampling approach," "An archaeological example: radiocarbon dating," --TR --CTR Gianluigi Pillonetto , Claudio Cobelli, Brief paper: Identifiability of the stochastic semi-blind deconvolution problem for a class of time-invariant linear systems, Automatica (Journal of IFAC), v.43 n.4, p.647-654, April, 2007 Xudong Jiang , Wee Ser, Online Fingerprint Template Improvement, IEEE Transactions on Pattern Analysis and Machine Intelligence, v.24 n.8, p.1121-1126, August 2002 Ranjan K. Dash , Erkki Somersalo , Marco E. Cabrera , Daniela Calvetti, An efficient deconvolution algorithm for estimating oxygen consumption during muscle activities, Computer Methods and Programs in Biomedicine, v.85 n.3, p.247-256, March, 2007

system identification;inverse problems;Markov chain Monte Carlo methods;dynamic systems;smoothing;bayesian estimation

297817

Bayesian Classification With Gaussian Processes.

AbstractWe consider the problem of assigning an input vector to one of m classes by predicting P(c|${\schmi x}$) for m. For a two-class problem, the probability of class one given ${\schmi x}$ is estimated by (y(${\schmi x}$)), where Gaussian process prior is placed on y(${\schmi x}$), and is combined with the training data to obtain predictions for new ${\schmi x}$ points. We provide a Bayesian treatment, integrating over uncertainty in y and in the parameters that control the Gaussian process prior; the necessary integration over y is carried out using Laplace's approximation. The method is generalized to multiclass problems (m > 2) using the softmax function. We demonstrate the effectiveness of the method on a number of datasets.

Introduction We consider the problem of assigning an input vector x to one out of m classes by predicting P (cjx) for classic example of this method is logistic regression. For a two-class problem, the probability of class 1 given x is estimated by oe(w T x+ b), where \Gammay ). However, this method is not at all "flexible", i.e. the discriminant surface is simply a hyperplane in x-space. This problem can be overcome, to some extent, by expanding the input x into a set of basis functions fOE(x)g, for example quadratic functions of the components of x. For a high-dimensional input space there will be a large number of basis functions, each one with an associated parameter, and one risks "overfitting" the training data. This motivates a Bayesian treatment of the problem, where the priors on the parameters encourage smoothness in the model. Putting priors on the parameters of the basis functions indirectly induces priors over the functions that can be produced by the model. However, it is possible (and we would argue, perhaps more natural) to put priors directly over the functions themselves. One advantage of function-space priors is that they can impose a general smoothness constraint without being tied to a limited number of basis functions. In the regression case where the task is to predict a real-valued output, it is possible to carry out non-parametric regression using Gaussian Processes (GPs); see, e.g. [25], [28]. The solution for the regression problem under a GP prior (and Gaussian noise model) is to place a kernel function on each training data point, with coefficients determined by solving a linear system. If the parameters ' that describe the Gaussian process are unknown, Bayesian inference can be carried out for them, as described in [28]. The Gaussian Process method can be extended to classification problems by defining a GP over y, the input to the sigmoid function. This idea has been used by a number of authors, although previous treatments typically do not take a fully Bayesian approach, ignoring uncertainty in both the posterior distribution of y given the data, and uncertainty in the parameters '. This paper attempts a fully Bayesian treatment of the problem, and also introduces a particular form of covariance function for the Gaussian process prior which, we believe, is useful from a modelling point of view. The structure of the remainder of the paper is as follows: Section 2 discusses the use of Gaussian processes for regression problems, as this is essential background for the classification case. In Section 3 we describe the application of Gaussian processes to two-class classification problems, and extend this to multiple-class problems in section 4. Experimental results are presented in section 5, followed by a discussion in section 6. This paper is a revised and expanded version of [1]. Gaussian Processes for regression It will be useful to first consider the regression problem, i.e. the prediction of a real valued output y for a new input value x , given a set of training data ng. This is of relevance because our strategy will be to transform the classification problem into a regression problem by dealing with the input values to the logistic transfer function. A stochastic process prior over functions allows us to specify, given a set of inputs, x the distribution over their corresponding outputs We denote this prior over functions as P (y), and similarly, P (y ; y) for the joint distribution including y . If we also specify P (tjy), the probability of observing the particular values actual values y (i.e. a noise model) then we have that Z Z Z Hence the predictive distribution for y is found from the marginalization of the product of the prior and the noise model. Note that in order to make predictions it is not necessary to deal directly with priors over function space, only n- or n + 1-dimensional joint densities. However, it is still not easy to carry out these calculations unless the densities involved have a special form. If P (tjy) and P (y ; y) are Gaussian then P (y jt) is a Gaussian whose mean and variance can be calculated using matrix computations involving matrices of size n \Theta n. Specifying P (y ; y) to be a multidimensional Gaussian (for all values of n and placements of the points x means that the prior over functions is a Gaussian process. More formally, a stochastic process is a collection of random variables fY (x)jx 2 Xg indexed by a set X . In our case X will be the input space with dimension d, the number of inputs. A GP is a stochastic process which can be fully specified by its mean function its covariance function C(x; x any finite set of Y -variables will have a joint multivariate Gaussian distribution. Below we consider GPs which have -(x) j 0. If we further assume that the noise model P (tjy) is Gaussian with mean zero and variance oe 2 I , then the predicted mean and variance at x are given by y e.g. [25]). 2.1 Parameterizing the covariance function There are many reasonable choices for the covariance function. Formally, we are required to specify functions which will generate a non-negative definite covariance matrix for any set of points From a modelling point of view we wish to specify covariances so that points with nearby inputs will give rise to similar predictions. We find that the following covariance function works well: l where x l is the lth component of x and is the vector of parameters that are needed to define the covariance function. Note that ' is analogous to the hyperparameters in a neural network. We define the parameters to be the log of the variables in equation (4) since these are positive scale-parameters. This covariance function can be obtained from a network of Gaussian radial basis functions in the limit of an infinite number of hidden units [27]. The w l parameters in equation 4 allow a different length scale on each input dimension. For irrelevant inputs, the corresponding w l will become small, and the model will ignore that input. This is closely related to the Automatic Relevance Determination (ARD) idea of MacKay [10] and Neal [15]. The v 0 variable specifies the overall scale of the prior. v 1 specifies the variance of a zero-mean offset which has a Gaussian distribution. The Gaussian process framework allows quite a wide variety of priors over functions. For example, the Ornstein-Uhlenbeck process (with covariance function C(x; x very rough sample paths which are not mean-square differentiable. On the other hand the squared exponential covariance function of equation 4 gives rise to an infinitely m.s. differentiable process. In general we believe that the GP method is a quite general-purpose route for imposing prior beliefs about the desired amount of smoothness. For reasonably high-dimensional problems, this needs to be combined with other modelling assumptions such as ARD. Another modelling assumption that may be used is to build up the covariance function as a sum of covariance functions, each one of which may depend on only some of the input variables (see section 3.3 for further details). 2.2 Dealing with parameters Given a covariance function it is straightforward to make predictions for new test points. However, in practical situations we are unlikely to know which covariance function to use. One option is to choose a parametric family of covariance functions (with a parameter vector ') and then either to estimate the parameters (for example, using the method of maximum likelihood) or to use a Bayesian approach where a posterior distribution over the parameters is obtained. These calculations are facilitated by the fact that the log likelihood l = log P (Dj') can be calculated analytically as log 2-; (5) where ~ Kj denotes the determinant of ~ K. It is also possible to express analytically the partial derivatives of the log likelihood with respect to the parameters @l ~ ~ (see, e.g. [11]). Given l and its derivatives with respect to ' it is straightforward to feed this information to an optimization package in order to obtain a local maximum of the likelihood. In general one may be concerned about making point estimates when the number of parameters is large relative to the number of data points, or if some of the parameters may be poorly determined, or if there may be local maxima in the likelihood surface. For these reasons the Bayesian approach of defining a prior s Figure 1: -(x) is obtained from y(x) by "squashing" it through the sigmoid function oe. distribution over the parameters and then obtaining a posterior distribution once the data D has been seen is attractive. To make a prediction for a new test point x one simply averages over the posterior distribution Z For GPs it is not possible to do this integration analytically in general, but numerical methods may be used. If ' is of sufficiently low dimension, then techniques involving grids in '-space can be used. If ' is high-dimensional it is very difficult to locate the regions of parameter-space which have high posterior density by gridding techniques or importance sampling. In this case Markov chain Monte Carlo methods may be used. These work by constructing a Markov chain whose equilibrium distribution is the desired distribution P ('jD); the integral in equation 7 is then approximated using samples from the Markov chain. Two standard methods for constructing MCMC methods are the Gibbs sampler and Metropolis-Hastings algorithms (see, e.g., [5]). However, the conditional parameter distributions are not amenable to Gibbs sampling if the covariance function has the form given by equation 4, and the Metropolis-Hastings algorithm does not utilize the derivative information that is available, which means that it tends to have an inefficient random-walk behaviour in parameter-space. Following the work of Neal [15] on Bayesian treatment of neural networks, Williams and Rasmussen [28] and Rasmussen [17] have used the Hybrid Monte Carlo (HMC) method of Duane et al [4] to obtain samples from P ('jD). The HMC algorithm is described in more detail in Appendix D. 3 Gaussian Processes for two-class classification For simplicity of exposition we will first present our method as applied to two-class problems; the extension to multiple classes is covered in section 4. By using the logistic transfer function to produce an output which can be interpreted as -(x), the probability of the input x belonging to class 1, the job of specifying a prior over functions - can be transformed into that of specifying a prior over the input to the transfer function, which we shall call the activation, and denote by y (see Figure 1). For the two-class problem we can use the logistic function \Gammay ). We will denote the probability and activation corresponding to input x i by - i and y i respectively. Fundamentally, the GP approaches to classification and regression problems are similar, except that the error model which is t - N(y; oe 2 ) in the regression case is replaced by t - Bern(oe(y)). The choice of v 0 in equation 4 will affect how "hard" the classification is; i.e. if -(x) hovers around 0:5 or takes on the extreme values of 0 and 1. Previous and related work to this approach is discussed in section 3.3. As in the regression case there are now two problems to address (a) making predictions with fixed parameters and (b) dealing with parameters. We shall discuss these issues in turn. 3.1 Making predictions with fixed parameters To make predictions when using fixed parameters we would like to compute - R requires us to find P (- for a new input x . This can be done by finding the distribution is the activation of - ) and then using the appropriate Jacobian to transform the distribution. Formally the equations for obtaining P (y jt) are identical to equations 1, 2, and 3. However, even if we use a GP prior so that P (y ; y) is Gaussian, the usual expression for classification data (where the t's take on values of 0 or 1), means that the marginalization to obtain P (y jt) is no longer analytically tractable. Faced with this problem there are two routes that we can follow: (i) to use an analytic approximation to the integral in equations 1-3 or (ii) to use Monte Carlo methods, specifically MCMC methods, to approximate it. Below we consider an analytic approximation based on Laplace's approximation; some other approximations are discussed in section 3.3. In Laplace's approximation, the integrand P (y ; yjt; ') is approximated by a Gaussian distribution centered at a maximum of this function with respect to y ; y with an inverse covariance matrix given by \Gammarr log P (y ; yjt; '). Finding a maximum can be carried out using the Newton-Raphson iterative method on y, which then allows the approximate distribution of y to be calculated. Details of the maximization procedure can be found in Appendix A. 3.2 Integration over the parameters To make predictions we integrate the predicted probabilities over the posterior P ('jt) / P (tj')P ('), as we saw in 2.2. For the regression problem P (tj') can be calculated exactly using P R P (tjy)P (yj')dy, but this integral is not analytically tractable for the classification problem. Let Using log log 2-: (8) By using Laplace's approximation about the maximum ~ y we find that log log 2-: We denote the right-hand side of this equation by log P a (tj') (where a stands for approximate). The integration over '-space also cannot be done analytically, and we employ a Markov Chain Monte Carlo method. Following Neal [15] and Williams and Rasmussen [28] we have used the Hybrid Monte Carlo (HMC) method of Duane et al [4] as described in Appendix D. We use log P a (tj') as an approximation for log P (tj'), and use broad Gaussian priors on the parameters. 3.3 Previous and related work Our work on Gaussian processes for regression and classification developed from the observation in [15] that a large class of neural network models converge to GPs in the limit of an infinite number of hidden units. The computational Bayesian treatment of GPs can be easier than for neural networks. In the regression case an infinite number of weights are effectively integrated out, and one ends up dealing only with the (hyper)parameters. Results from [17] show that Gaussian processes for regression are comparable in performance to other state-of-the-art methods. Non-parametric methods for classification problems can be seen to arise from the combination of two different strands of work. Starting from linear regression, McCullagh and Nelder [12] developed generalized linear models (GLMs). In the two-class classification context, this gives rise to logistic regression. The other strand of work was the the development of non-parametric smoothing for the regression problem. Viewed as a Gaussian process prior over functions this can be traced back at least as far as the work of Kolmogorov and Wiener in the 1940s. Gaussian process prediction is well known in the geostatistics field (see, e.g. [3]) where it is known as "kriging". Alternatively, by considering "roughness penalties" on functions, one can obtain spline methods; for recent overviews, see [25] and [8]. There is a close connection between the GP and roughness penalty views, as explored in [9]. By combining GLMs with non-parametric regression one obtains what we shall call a non-parametric GLM method for classification. Early references to this method include [21] and [16], and discussions can also be found in texts such as [8] and [25]. There are two differences between the non-parametric GLM method as it is usually described and a Bayesian treatment. Firstly, for fixed parameters the non-parametric GLM method ignores the uncertainty in y and hence the need to integrate over this (as described in section 3.1). The second difference relates to the treatment of the parameters '. As discussed in section 2.2, given parameters ', one can either attempt to obtain a point estimate for the parameters or to carry out an integration over the posterior. Point estimates may be obtained by maximum likelihood estimation of ', or by cross-validation or generalized cross-validation (GCV) methods, see e.g. [25, 8]. One problem with CV-type methods is that if the dimension of ' is large, then it can be computationally intensive to search over a region/grid in parameter-space looking for the parameters that maximize the criterion. In a sense the HMC method described above are doing a similar search, but using gradient information 1 , and carrying out averaging over the posterior distribution of parameters. In defence of (G)CV methods, we note Wahba's comments (e.g. in [26], referring back to [24]) that these methods may be more robust against an unrealistic prior. One other difference between the kinds of non-parametric GLM models usually considered and our method is the exact nature of the prior that is used. Often the roughness penalties used are expressed in terms of a penalty on the kth derivative of y(x), which gives rise to a power law power spectrum for the prior on y(x). There can also be differences over parameterization of the covariance function; for example it is unusual to find parameters like those for ARD introduced in equation 4 in non-parametric GLM models. On the other hand, Wahba et al [26] have considered a smoothing spline analysis of variance (SS-ANOVA) decomposition. In Gaussian process terms, this builds up a prior on y as a sum of priors on each of the functions in the decomposition ff y ff The important point is that functions involving all orders of interaction (from univariate functions, which on their own give rise to an additive model) are included in this sum, up to the full interaction term which is the only one that we are using. From a Bayesian point of view questions as to the kinds of priors that are appropriate is an interesting modelling issue. There has also been some recent work which is related to the method presented in this paper. In section 3.1 we mentioned that it is necessary to approximate the integral in equations 1-3 and described the use of Laplace's approximation. Following the preliminary version of this paper presented in [1], Gibbs and MacKay [7] developed an alternative analytic approximation, by using variational methods to find approximating Gaussian distributions that bound the marginal likelihood P (tj') above and below. These approximate distributions are then used to predict P (y jt; ') and thus - -(x ). For the parameters, Gibbs and MacKay estimated ' by maximizing their lower bound on P (tj'). It is also possible to use a fully MCMC treatment of the classification problem, as discussed in the recent paper of Neal [14]. His method carries out the integrations over the posterior distributions of y and ' simultaneously. It works by generating samples from P (y; 'jD) in a two stage process. Firstly, for fixed ', each of the n individual y i 's are updated sequentially using Gibbs sampling. This ``sweep'' takes time O(n 2 ) once the matrix K \Gamma1 has been computed (in time O(n 3 )), so it actually makes sense to perform quite a few Gibbs sampling scans between each update of the parameters, as this probably makes the Markov chain mix faster. Secondly, the parameters are updated using the Hybrid Monte Carlo method. To make predictions, one averages over the predictions made by each It would be possible to obtain derivatives of the CV-score with respect to ', but this has not, to our knowledge, been used in practice. 4 GPs for multiple-class classification The extension of the preceding framework to multiple classes is essentially straightforward, although notationally more complex. Throughout we employ a one-of-m class coding scheme 2 , and use the multi-class analogue of the logistic function-the softmax function-to describe the class probabilities. The probability that an instance labelled by i is in class c is denoted by - i c , so that an upper index to denotes the example number, and a lower index the class label. Similarly, the activations associated with the probabilities are denoted by y i c . Formally, the link function relates the activations and probabilities through c which automatically enforces the constraint 1. The targets are similarly represented by t i c , and are specified using a one-of-m coding. The log likelihood takes the form c , which for the softmax link function gives c As for the two class case, we shall assume that the GP prior operates in activation space; that is we specify the correlations between the activations y i c . One important assumption we make is that our prior knowledge is restricted to correlations between the activations of a particular class. Whilst there is no difficulty in extending the framework to include inter-class correlations, we have not yet encountered a situation where we felt able to specify such correlations. Formally, the activation correlations take the form, hy i c (12) where K i;i 0 c is the element of the covariance matrix for the cth class. Each individual correlation matrix K c has the form given by equation 4 for the two-class case. We shall use a separate set of parameters for each class. The use of m independent processes to perform the classification is redundant, but forcing the activations of one process to be (say) zero would introduce an arbitrary asymmetry into the prior. For simplicity, we introduce the augmented vector notation, where, as in the two-class case, y c denotes the activation corresponding to input x for class c; this notation is also used to define t + and -+ . In a similar manner, we define y, t and - by excluding the values corresponding to the test point x . Let y With this definition of the augmented vectors, the GP prior takes the form, ae oe where, from equation 12, the covariance matrix K + is block diagonal in the matrices, K m . Each individual matrix K c expresses the correlations of activations within class c. As in the two-class case, to use Laplace's approximation we need to find the mode of P (y jt). The procedure is described in Appendix C. As for the two-class case, we make predictions for -(x ) by averaging the softmax function over the Gaussian approximation to the posterior distribution of y . At present, we simply estimate this integral using 1000 draws from a Gaussian random vector generator. That is, the class is represented by a vector of length m with zero entries everywhere except for the correct component which contains 1. 5 Experimental results When using the Newton-Raphson algorithm, - was initialized each time with entries 1=m, and iterated until the mean relative difference of the elements of W between consecutive iterations was less than 10 \Gamma4 . For the HMC algorithm, the same step size " is used for all parameters, and should be as large as possible while keeping the rejection rate low. We have used a trajectory made up of leapfrog steps, which gave a low correlation between successive states. The priors over parameters were set to be Gaussian with a mean of \Gamma3 and a standard deviation of 3. In all our simulations a step size produced a low rejection rate (! 5%). The parameters corresponding to the w l 's were initialized to \Gamma2 and that for v 0 to 0. The sampling procedure was run for 200 iterations, and the first third of the run was discarded; this "burn-in" is intended to give the parameters time to come close to their equilibrium distribution. Tests carried out using the R-CODA package 3 on the examples in section 5.1 suggested that this was indeed effective in removing the transients, although we note that it is widely recognized (see, e.g. [2]) that determining when the equilibrium distribution has been reached is a difficult problem. Although the number of iterations used is much less than typically used for MCMC methods it should be remembered that (i) each iteration involves leapfrog steps and (ii) that by using HMC we aim to reduce the "random walk" behaviour seen in methods such as the Metropolis algorithm. Autocorrelation analysis for each parameter indicated, in general, that low correlation was obtained after a lag of a few iterations. The MATLAB code which we used to run our experiments is available from ftp://cs.aston.ac.uk/neural/willicki/gpclass/. 5.1 Two classes We have tried out our method on two well known two class classification problems, the Leptograpsus crabs and Pima Indian diabetes datasets 4 . We first rescale the inputs so that they have mean of zero and unit variance on the training set. Our Matlab implementations for the HMC simulations for both tasks each take several hours on a SGI Challenge machine (200MHz R10000), although good results can be obtained in much less time. We also tried a standard Metropolis MCMC algorithm for the Crabs problem, and found similar results, although the sampling by this method is somewhat slower than that for HMC. The results for the Crabs and Pima tasks, together with comparisons with other methods (from [20] and [18]) are given in Tables 1 and 2 respectively. The tables also include results obtained for Gaussian processes using (a) estimation of the parameters by maximizing the penalised likelihood (found using 20 iterations of a scaled conjugate gradient optimiser) and (b) Neal's MCMC method. Details of the set-up used for Neal's method are given in Appendix E. In the Leptograpsus crabs problem we attempt to classify the sex of crabs on the basis of five anatomical attributes, with an optional additional colour attribute. There are 50 examples available for crabs of each sex and colour, making a total of 200 labelled examples. These are split into a training set of 20 crabs of each sex and colour, making 80 training examples, with the other 120 examples used as the test set. The performance of our GP method is equal to the best of the other methods reported in [20], namely a 2 hidden unit neural network with direct input to output connections, a logistic output unit and trained with maximum likelihood (Network(1) in Table 1). Neal's method gave a very similar level of performance. We also found that estimating the parameters using maximum penalised likelihood (MPL) gave similar performance with less than a minute of computing time. For the Pima Indians diabetes problem we have used the data as made available by Prof. Ripley, with his training/test split of 200 and 332 examples respectively [18]. The baseline error obtained by simply classifying each record as coming from a diabetic gives rise to an error of 33%. Again, ours and Neal's GP methods are comparable with the best alternative performance, with an error of around 20%. It is encouraging that the results obtained using Laplace's approximation and Neal's method are similar 5 . We also estimated the parameters using maximum penalised likelihood, rather than Monte Carlo integration. The performance in this case was a little worse, with 21.7% error, but for only 2 minutes computing time. 3 Available from the Comprehensive R Archive Network at http://www.ci.tuwien.ac.at. 4 Available from http://markov.stats.ox.ac.uk/pub/PRNN. 5 The performance obtained by Gibbs and MacKay in [7] was similar. Their method made 4 errors in the crab task (with colour given), and 70 errors on the Pima dataset. Method Colour given Colour not given Neural Network(1) 3 3 Neural Network(2) 5 3 Linear Discriminant 8 8 Logistic regression 4 4 PP regression (4 ridge Gaussian Process (Laplace 3 3 Approximation, HMC) Gaussian Process (Laplace 4 3 Approximation, MPL) Gaussian Process (Neal's method) 4 3 Table 1: Number of test errors for the Leptograpsus crabs task. Comparisons are taken from from Ripley (1996) and Ripley (1994) respectively. Network(2) used two hidden units and the predictive approach (Ripley, 1993) which uses Laplace's approximation to weight each network local minimum. Method Pima Indian diabetes Neural Network 75+ Linear Discriminant 67 Logistic Regression 66 PP regression (4 ridge functions) 75 Gaussian Mixture 64 Gaussian Process (Laplace 68 Approximation, HMC) Gaussian Process (Laplace 69 Approximation, MPL) Gaussian Process (Neal's method) 68 Table 2: Number of test errors on the Pima Indian diabetes task. Comparisons are taken from from Ripley (1996) and Ripley (1994) respectively. The neural network had one hidden unit and was trained with maximum likelihood; the results were worse for nets with two or more hidden units (Ripley, 1996). Analysis of the posterior distribution of the w parameters in the covariance function (equation be informative. Figure 5.1 plots the posterior marginal mean and 1 standard deviation error bars for each of the seven input dimensions. Recalling that the variables are scaled to have zero mean and unit variance, it would appear that variables 1 and 3 have the shortest lengthscales (and therefore the most variability) associated with them. 5.2 Multiple classes Due to the rather long time taken to run our code, we were only able to test it on relatively small problems, by which we mean only a few hundred data points and several classes. Furthermore, we found that a full Bayesian integration over possible parameter settings was beyond our computational means, and we therefore had to be satisfied with a maximum penalised likelihood approach. Rather than using the potential and its gradient in a HMC routine, we now simply used them as inputs to a scaled conjugate gradient optimiser (based on [13]) instead, attempting to find a mode of the class posterior, rather than to average over the posterior distribution. We tested the multiple class method on the Forensic Glass dataset described in [18]. This is a dataset of 214 examples with 9 inputs and 6 output classes. Because the dataset is so small, the performance is Figure 2: Plot of the log w parameters for the Pima dataset. The circle indicates the posterior marginal mean obtained from the HMC run (after burn-in), with one standard deviation error bars. The square symbol shows the log w-parameter values found by maximizing the penalized likelihood. The variables are 1. the number of pregnancies, 2. plasma glucose concentration, 3. diastolic blood pressure, 4. triceps skin fold thickness, 5. body mass index, 6. diabetes pedigree function, 7. age. For comparison, Wahba et al (1995) using generalized linear regression, found that variables 1, 2 5 and 6 were the most important. estimated from using 10-fold cross validation. Computing the penalised maximum likelihood estimate of our multiple GP method took approximately 24 hours on our SGI Challenge and gave a classification error rate of 23.3%. As we see from Table 3, this is comparable to the best of the other methods. The performance of Neal's method is surprisingly poor; this may be due to the fact that we allow separate parameters for each of the y processes, while these are constrained to be equal in Neal's code. There are also small but perhaps significant differences in the specification of the prior (see Appendix E for details). 6 Discussion In this paper we have extended the work of Williams and Rasmussen [28] to classification problems, and have demonstrated that it performs well on the datasets we have tried. We believe that the kinds of Gaussian Method Forensic Glass Neural Network (4HU) 23.8% Linear Discriminant 36% PP regression (5 ridge functions) 35% Gaussian Mixture 30.8% Decision Tree 32.2% Gaussian Process (LA, MPL) 23.3% Gaussian Process (Neal's method) 31.8% Table 3: Percentage of test error for the Forensic Glass problem. See Ripley (1996) for details of the methods. process prior we have used are more easily interpretable than models (such as neural networks) in which the priors are on the parameterization of the function space. For example, the posterior distribution of the ARD parameters (as illustrated in Figure 5.1 for the Pima Indians diabetes problem) indicates the relative importance of various inputs. This interpretability should also facilitate the incorporation of prior knowledge into new problems. There are quite strong similarities between GP classifiers and support-vector machines (SVMs) [23]. The SVM uses a covariance kernel, but differs from the GP approach by using a different data fit term (the maximum margin), so that the optimal y is found using quadratic programming. The comparison of these two algorithms is an interesting direction for future research. A problem with methods based on GPs is that they require computations (trace, determinants and linear solutions) involving n \Theta n matrices, where n is the number of training examples, and hence run into problems on large datasets. We have looked into methods using Bayesian numerical techniques to calculate the trace and determinant [22, 6], although we found that these techniques did not work well for the (relatively) small size problems on which we tested our methods. Computational methods used to speed up the quadratic programming problem for SVMs may also be useful for the GP classifier problem. We are also investigating the use of different covariance functions and improvements on the approximations employed. Acknowledgements We thank Prof. B. Ripley for making available the Leptograpsus crabs, Pima Indian diabetes and Forensic Glass datasets. This work was partially supported by EPSRC grant GR/J75425, Novel Developments in Learning Theory for Neural Networks, and much of the work was carried out at Aston University. The authors gratefully acknowledge the hospitality provided by the Isaac Newton Institute for Mathematical Sciences (Cambridge, UK) where this paper was written up. We thank Mark Gibbs, David MacKay and Radford Neal for helpful discussions, and the anonymous referees for their comments which helped improve the paper. Appendix Maximizing case We describe how to find iteratively the vector y + so that P (y + jt) is maximized. This material is also covered in [8] x5.3.3 and [25] x9.2. We provide it here for completeness and so that the terms in equation 9 are well-defined. the complete set of activations. By Bayes' theorem log log As P (t) does not depend on y + (it is just a normalizing factor), the maximum of P (y + jt) is found by maximizing \Psi + with respect to y + . Using log log 2- (14) where K+ is the covariance matrix of the GP evaluated at x . \Psi is defined similarly in equation 8. K+ can be partitioned in terms of an n \Theta n matrix K, a n \Theta 1 vector k and a scalar k , viz. As y only enters into equation 14 in the quadratic prior term and has no data point associated with it, maximizing with respect to y + can be achieved by first maximizing \Psi with respect to y and then doing the further quadratic optimization to determine y . To find a maximum of \Psi we use the Newton-Raphson iteration y new Differentiating equation 8 with respect to y we find where the 'noise' matrix is given by This results in the iterative equation, To avoid unnecessary inversions, it is usually more convenient to rewrite this in the form Note that \Gammarr\Psi is always positive definite, so that the optimization problem is convex. Given a converged solution ~ y for y, y can easily be found using y -), as is the W with a zero appended in the (n diagonal position. Given the mean and variance of y it is then easy to find R the mean of the distribution of P (- jt). In order to calculate the Gaussian integral over the logistic sigmoid function, we employ an approximation based on the expansion of the sigmoid function in terms of the error function. As the Gaussian integral of an error function is another error function, this approximation is fast to compute. Specifically, we use a basis set of five scaled error functions to interpolate the logistic sigmoid at chosen points 6 . This gives an accurate approximation (to to the desired integral with a small computational cost. The justification of Laplace's approximation in our case is somewhat different from the argument usually put forward, e.g. for asymptotic normality of the maximum likelihood estimator for a model with a finite number of parameters. This is because the dimension of the problem grows with the number of data points. However, if we consider the "infill asymptotics" (see, e.g. [3]), where the number of data points in a bounded region increases, then a local average of the training data at any point x will provide a tightly localized estimate for -(x) and hence y(x) (this reasoning parallels more formal arguments found in [29]). Thus we would expect the distribution P (y) to become more Gaussian with increasing data. Appendix B: Derivatives of log P a (tj') wrt '. For both the HMC and MPL methods we require the derivative of l a log P a (tj') with respect to components of ', for example ' k . This derivative will involve two terms, one due to explicit dependencies of l a = log 2- on ' k , and also because a change in ' will cause a change in ~ y. However, as ~ y is chosen so that r\Psi(y)j y= ~ @l a \Gamma2 @ log jK The dependence of jK \Gamma1 +W j on ~ y arises through the dependence of W on ~ -, and hence ~ y. By differentiating ~ -), one obtains @~ y and hence the required derivative can be calculated. Appendix Maximizing Multiple-class case The GP prior and likelihood, defined by equations 13 and 11, define the posterior distribution of activations, jt). As in Appendix A we are interested in a Laplace approximation to this posterior, and therefore need to find the mode with respect to y Dropping unnecessary constants, the multi-class analogue of equation 14 is c exp y i 6 In detail, we used the basis functions erf(-x)) for These were used to interpolate oe(x) at By the same principle as in Appendix A, we define \Psi by analogy with equation 8, and first optimize \Psi with respect to y, afterwards performing the quadratic optimization of \Psi + with respect to y . In order to optimize \Psi with respect to y, we make use of the Hessian given by where K is the mn \Theta mn block-diagonal matrix with blocks K c , m. Although this is in the same form as for the two class case, equation 17, there is a slight change in the definition of the 'noise' matrix, W . A convenient way to define W is by introducing the matrix \Pi which is a mn \Theta n matrix of the form Using this notation, we can write the noise matrix in the form of a diagonal matrix and an outer product, As in the two-class case, we note that \Gammarr\Psi is again positive definite, so that the optimization problem is convex. The update equation for iterative optimization of \Psi with respect to the activations y then follows the same form as that given by equation 18. The advantage of the representation of the noise matrix in equation 23 is that we can then invert matrices and find their determinants using the identities, and As A is block-diagonal, it can be inverted blockwise. Thus, rather than requiring determinants and inverses of a mn \Theta mn matrix, we only need to carry out expensive matrix computations on n \Theta n matrices. The resulting update equations for y are then of the same form as given in equation 18, where the noise matrix and covariance matrices are now in their multiple class form. Essentially, these are all the results needed to generalize the method to the multiple-class problem. Although, as we mentioned above, the time complexity of the problem does not scale with the m 3 , but rather m (due to the identities in equations 24, 25), calculating the function and its gradient is still rather expensive. We have experimented with several methods of mode finding for the Laplace approximation. The advantage of the Newton iteration method is its fast quadratic convergence. An integral part of each Newton step is the calculation of the inverse of a matrix M acting upon a vector, i.e. M \Gamma1 b . In order to speed up this particular step, we used a conjugate gradient (CG) method to solve iteratively the corresponding linear system b. As we repeatedly need to solve the system (because W changes as y is updated), it saves time not to run the CG method to convergence each time it is called. In our experiments the CG algorithm was terminated when 1=n The calculation of the derivative of log P a (tj') wrt ' in the multiple-class case is analogous to the two-class case described in Appendix B. Appendix D: Hybrid Monte Carlo HMC works by creating a fictitious dynamical system in which the parameters are regarded as position variables, and augmenting these with momentum variables p. The purpose of the dynamical system is to give the parameters "inertia" so that random-walk behaviour in '-space can be avoided. The total energy, H , of the system is the sum of the kinetic energy, and the potential energy, E. The potential energy is defined such that p('jD) / exp(\GammaE), i.e. We sample from the joint distribution for ' and p given by P ('; p) / exp(\GammaE \Gamma K); the marginal of this distribution for ' is the required posterior. A sample of parameters from the posterior can therefore be obtained by simply ignoring the momenta. Sampling from the joint distribution is achieved by two steps: (i) finding new points in phase space with near-identical energies H by simulating the dynamical system using a discretised approximation to Hamiltonian dynamics, and (ii) changing the energy H by Gibbs sampling the momentum variables. Hamilton's first order differential equations for H are approximated using the leapfrog method which requires the derivatives of E with respect to '. Given a Gaussian prior on ', log P (') is straightforward to differentiate. The derivative of log P a (tj') is also straightforward, although implicit dependencies of ~ y (and hence ~ -) on ' must be taken into account as described in Appendix B. The calculation of the energy can be quite expensive as for each new ', we need to perform the maximization required for Laplace's approximation, equation 9. This proposed state is then accepted or rejected using the Metropolis rule depending on the final energy H (which is not necessarily equal to the initial energy H because of the discretization). Appendix E: Simulation set-up for Neal's code We used the fbm software available from http://www.cs.utoronto.ca/~radford/fbm.software.html. For example, the commands used to run the Pima example are model-spec pima1.log binary gp-gen pima1.log fix 0.5 1 mc-spec pima1.log repeat 4 scan-values 200 heatbath hybrid 6 0.5 gp-mc pima1.log 500 which follow closely the example given in Neal's documentation. The gp-spec command specifies the form of the Gaussian process, and in particular the priors on the parameters v 0 and the w's (see equation 4). The expression 0.05:0.5 specifies a Gamma-distribution prior on v 0 , and x0.2:0.5:1 specifies a hierarchical Gamma prior on the w's. Note that a ``jitter'' of 0:1 is also specified on the prior covariance function; this improves conditioning of the covariance matrix. The mc-spec command gives details of the MCMC updating procedure. It specifies 4 repetitions of 200 scans of the y values followed by 6 HMC updates of the parameters (using a step-size adjustment factor of 0.5). gp-mc specifies that this is sequence is carried out 500 times. We aimed for a rejection rate of around 5%. If this was exceeded, the stepsize reduction factor was reduced and the simulation run again. --R Statistics for Spatial Data. Hybrid Monte Carlo. Bayesian Data Analysis. Efficient Implementation of Gaussian Processes. Variational Gaussian Process Classifiers. 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gaussian processes;Markov chain Monte Carlo;classification problems;hybrid Monte Carlo;bayesian classification;parameter uncertainty

297877

A Positive Acknowledgment Protocol for Causal Broadcasting.

AbstractCausal broadcasting has been introduced to reduce the asynchrony of communication channels inside groups of processes. It states that if two broadcast messages are causally related by the happened-before relation, these messages are delivered in their sending order to each process of the group. Even though protocols implementing causal broadcasting do not add control messages, they suffer from the typical pitfall of the timestamping technique: To ensure causal ordering, application messages have to piggyback a vector time of counters whose range of variation is unbounded. In this paper, we investigate such a range and define the concept of causal window of a process in which all counters of a vector time of a just arrived message at that process fall. We prove that, by using a causal broadcasting (one-to-all) protocol that follows a positive acknowledgment method, the width of the causal window of each process is limited. This allows a modulo k implementation of vector times when considering k greater than the width of the causal window of each process. The protocol is applicable to data link or transport layers using acknowledge messages to ensure reliable transfer of data. The paper also proposes two variants of the protocol based on causal windows. Both of them increase the concurrency of the protocol at the expense of wider causal windows.

Introduction Asynchrony of communication channels is one of the major causes of nondeterminism in distributed systems. The concept of causal ordering of messages has been introduced in the context of broadcasting communication by Birman and Joseph [8] in order to reduce such an asynchrony. Causal ordering means if two broadcast messages are causally related [14], they are delivered in their sending order to each process. In light of this, when a message is delivered to a process, all messages that causally precede it have been already delivered to that process. To master asynchrony, other communication modes have been defined such as FIFO, Rendezvous and logical instantaneous ordering [20]. From the user viewpoint, causal ordering increases the control of a distributed application compared to a simple FIFO ordering, at the cost of a reduction of the potential concurrency of the distributed application. Compared with logically instantaneous communication, causal ordering provides more concurrency and simplicity of implementation. Moreover, causal ordering is not prone to deadlock as Rendezvous, being an asynchronous paradigm of communication. Actually, causal ordering extends the concept of FIFO channels connecting one sender and one receiver to systems connecting several senders and one receiver. Causal ordering has been proved to be very useful in taking snapshots of distributed applications, in controlling distributed applications, in managing replicated data and in allowing consistent observations of distributed computations [10, 18]. Recently, extensions of causal ordering have been proposed to cope with mobile computing environments [16] and with unreliable networks and distributed applications whose messages have limited time validity [5, 6]. Moreover, the concept of causal ordering is not limited to message-passing environments. In the context of shared-memory systems, a causal memory has been introduced by Ahamad et al. in [1] as a consistency criterion. Such a criterion does not introduce latencies while executing read and write operations. Even though several interesting protocols implementing causal ordering appeared in the literature [8, 9, 18, 21], this communication mode is not yet widely used in commercial platforms because such protocols suffer from the typical pitfall of the timestamping (logical or physical) technique: to ensure causal ordering, in the context of broadcasting, application messages have to piggyback a vector time of unbounded integers (counters) whose size is given by the number of processes [15], this vector represents actually the control information of a protocol. However, data-link and transport layers of communication systems use messages, called acknowledgments (acks, for short), to indicate the successful reception of data. Ack messages, produced by such layers, are actually a source of information about the causal relations among application messages that could be used to reduce the amount of control information of causal ordering pro- tocols. In this paper we introduce the notion of causal window and propose a causal broadcasting protocol which exploits the implicit information provided by ack messages 1 . A causal window of a process represents the range of variation in which all counters of a vector time of a just arrived message at that process fall. We prove that, by using a causal broadcasting protocol that follows a positive acknowledgment (PAK) method [17, 23], the width of the causal window of each process is bounded. This allows a modulo k implementation of vector times when considering k greater than the width of the causal window of each process. We first propose a PAK causal broadcasting protocol in which a process can send a message only when acks of the previous message, sent by the same process, have been received. We analyze then the general case in which a credit ct - 1 is associated with each sending process; in this case, a process can send ct consecutive messages before receiving the corresponding acks. Finally, we investigate the case in which a process employs a positive/negative (PAK/NAK) scheme, i.e., a process can send a sequence of ct messages without being acknowledged and then an ack message is required from the other processes after the receipt of the ct-th message sent by the same process. Credits and the use of PAK/NAK scheme allow an increase of the concurrency of the protocol (decreasing the number of the internal protocol synchronizations), but enlarge the dimension of the causal window. The protocol we propose could be employed as a part of the flow control of a transport layer 2 providing causal communication to the above layer. For example, current group communication systems (e.g., ISIS [10]) implement causal protocols on the top of a FIFO flow control by piggybacking on each application message a vector of unbounded integer. The remainder of this paper is organized as follows. In Section 2 the general model of a distributed computation, the concept of causal relation among events, vector times and the causal ordering communication mode are introduced. Section 3 presents the causal window notion. Section 4 shows a causal broadcasting protocol based on causal windows when considering the credit of the sender equal to one. At this end, in the same section, the positive acknowledgment method and the modulo k implementation of vector times are introduced. Section 5 proposes the two variations of the protocol of Section 4 based on a credit and a PAK/NAK scheme respectively. 1 Some interesting algorithms that exploit the implicit information provided by ack messages to guarantee FIFO and reliable channels can be found in [7, 11, 22]. 2 Examples of transport layers that use ack messages for data transferring are, among others, TCP, OSI/TP4, VTMP and Delta-t [13]. 2 Model of Distributed Computations 2.1 Distributed System A distributed system is a finite set P of n processes fP that communicate only by broadcasting messages 3 . The underlying system, where processes execute, is composed of n processors (for simplicity's sake, we assume one process per processor) that can exchange messages. We assume that each pair of processes is connected by a reliable 4 , asynchronous and FIFO logical channel (transmission delays are unpredictable). Processors do not have a shared memory and there is no bound for their relative speeds. 2.2 Distributed executions Execution of a process P i produces a sequence of events which can be classified as: broadcast (bcast) events, deliver (dlv) events and internal events. An internal event may change only local variables, broadcast or delivery events involve communication. In particular, each broadcast event produces delivery events, one for each process. Let a and b be two events occurred in a process P i , a precedes b in P i , denoted aOE i b, iff a has been produced before b. Let m be a message and a and b two events, a precedes b, denoted a OE m b, iff a is the bcast(m) event and b is the dlv(m) event. A distributed computation can be represented as a partial order of events b E is the set of all events and ! is the happened-before relation [14]. This relation is the transitive closure of the union of OE i n) and OE m , it is denoted by !, i.e., Hereafter, we call M( b E) the set of all messages exchanged in b E and we do not consider internal events as they do not affect the interprocess ordering of events. Let us, finally, introduce the notion of the immediate predecessor of a message m. Definition 2.1 A message m 1 is an immediate predecessor of a message paper, we consider a message the atomic unit of data movement in the system. Results of the following Sections apply even though we consider packets or byte streams as atomic data unit. 4 A detailed description of the protocol proposed in Section 4 in the case of unreliable channel is out of the aims of this paper. The interested reader can refer to [3] for such a description. E); It is to be noted that a message m can have n immediate predecessors one for each process. As an example the message m 1 depicted in Figure 4 is not an immediate predecessor of m 2 due to message m x . 2.3 Vector Times To capture the causality relation between relevant events of a distributed computation, vector times were introduced simultaneously and independently by Fidge [12] and Mattern [15]. A vector time for a process P i , denoted V T i , is a vector of counters whose dimension is equal to the number of knowledge of the number of relevant events produced by P j . Each relevant event a is associated with a vector time (V T a ) and a process P i updates its vector time according to the following rules: 1. When P i starts its execution, each component of V T i is initialized to zero; 2. When a relevant event is produced by 3. When a message m is sent by A copy of V T i is piggybacked on message m (denoted V Tm ); 4. When a message m, sent by P j , arrives at P i , it updates its vector time in the following way: Let a and b be two relevant events and V T a and V T b the vector times associated with, according to properties of vector times [12, 15], we have: when considering V T a As we are interested in a broadcast environment, in the following, we assume broadcast events as the only relevant events, actually P i 's knowledge of the number of messages broadcast by P j . In this particular setting, and when considering causal communication, the above protocol can be simplified as shown in the next section. 2.4 Causal Ordering Causal ordering states that the order in which messages are delivered to the application cannot violate the happened-before relation of the corresponding broadcast events [8]. More formally, Definition 2.2 A distributed computation b respects causal ordering if for any two broadcast messages E) we have: ng :: dlv(m 1 A first implementation of such an abstraction has been embedded in the ISIS system [8, 10]. It consists of adding a protocol over a reliable underlying system such that events of a distributed computation be causally ordered at the process level 5 . At this end, deliveries are done by delaying, by means of a delivery condition, those messages arrived too early at the underlying system. A simple broadcast protocol, similar to the one presented by Birman et al. in [9], is shown in Figure 1. It shows the behavior of process P i when sending and upon the arrival of a message. The algorithm includes some vector time management rules (line S1,S2,S3 and R2) plus a delivery condition DC(m) (line R1) associated with a message m. A message m is delivered to a process P i as soon as the vector time, it carries (V Tm ) does not contain knowledge of messages sent to, but not delivered by, process P i 6 . Formally, ng 3 The Causal Window In this section we investigate the range of variation of the values stored in the counters of vector times during the evaluation of the delivery condition of a generic destination process. Upon the arrival of a message m at process P i , the value of vector time counters, V Tm and V T i , involved in the delivery condition DC(m) generate three cases: 1. V Tm 1. There are consecutive messages, sent by process P j , that causally precede m and that have not arrived at P i as shown in Figure 2.a. Message m, if delivered, violates causal ordering. 5 Other interesting point-to-point implementations of causal ordering can be found in [18, 21].This fact makes a part of the rule 3 of Section 2.3 (i.e., (8h 6= useless in a causal broadcasting protocol. init for each h 2 procedure BCAST(m; is the message, P i is the sender % begin for each h 2 ng do send (m; V Tm) to P h ; od % event bcast(m) % (S2) end. when (m;V Tm) arrives at P i from begin ng event dlv(m) % (R3) end. Figure 1: A simple causal broadcasting protocol 2. 1. There are CO j consecutive messages, sent by process P j and delivered to P i , that are concurrent to message m, as shown in Figure 2.b. Message m can be delivered to P i without violating causal ordering. 3. sent by P j and delivered to P i , is an immediate predecessor of m as shown in Figure 2.c. Message m can then be delivered to P i without violating causal ordering. (a) (b) CO j messages messages (c) Figure 2: Values of vector times in the delivery condition. Figure 3: The Causal Window CW j If all counters of V Tm fall either in the case 3 or in the case 2, message m can be delivered. Hence, upon the arrival of message m at process P i , counters of V Tm fall in a range of variation that spans between V T A causal window CW i is composed of a set of windows CW j i one for each process P j . The number i represents the width of the window CW j Figure 3). To implement a causal broadcasting protocol employing modulo k vector times, we have to show the boundedness of the width of the causal windows. In the general setting, as the one described in the previous section, where transmission times are unpredictable, the width of CW i are non-limited. 4 A PAK Protocol based on Causal Windows 4.1 Positive Acknowledgement Method To get a limited causal window, we assume processes follow a stop-and-wait approach. A process broadcasts a message ( bcast(m) event) and waits for all the acks (n a:ack(m) events) before executing any other broadcast event. Once all acks have arrived, such a broadcast message is said to be "fully acknowledged" (f:ack(m) event). On the other side, each time a process receives a broadcast message (arr(m) and then dlv(m) events), it sends an ack (s:ack(m)). Hence, at the underlying system level six types of events occur and only bcast and dlv events are visible to the application. So the processing of a broadcast message m, sent by P i , produces the following poset of events, denoted PO(m): OEm arr(m) OE1 dlv(m) OE1 s:ack(m) OEm OEm arr(m) OEn dlv(m) OEn s:ack(m) OEm the stop-and-wait approach implies a send condition SC 1 between any two successive messages sent by the same process P i . Formally, This synchronization is local and can be easily implemented by a boolean variable processing broadcast (initialised to FALSE) in each process. The value TRUE indicates that the process has broadcast a message m and it is waiting for f:ack(m) event. As soon as the event occurs processing broadcast toggles enabling other broadcast of messages. A remark on group communication. In a group communication system (e.g. ISIS [10], TRANSIS [2] etc.) the occurrence of the event fully:ack of a message m corresponds to the notion of stability of that message [4, 9], i.e., the sender of m learns that all the members of the group have delivered m. In fact the use of ack messages is one of the methods to diffuse stability information in a group of processes (other methods employ "gossiping" and piggybacking). The notion of stability is a key point in many group communication problems such as security [19] and large scale settings [4] just to name a few. So, informally, the condition SC 1 can be restated as follows: a process cannot broadcast a message in a group till the previous one, it sent, is declared stable. This condition is very conservative and implies a strong synchronization between each pair of successive messages sent by the same process. In Section 5 we present two variations that weaken that synchronization. 4.2 Causal Windows with Limited Width In this subsection we prove that, the causal window of a PAK causal broadcasting protocol based on the send condition SC 1 is limited: Lemma 4.1 Let b E be a distributed computation and m 1 and E) be messages such that protocol based on SC 1 ensures if there is a causal ordering violation between m 1 and m 2 then Proof (by contradiction) As shown in Figure 4, suppose there is a causal ordering violation between two messages m 1 and m 2 sent by P i and P j (i 6= respectively (i.e., bcast(m 1 that is, there exists a message m x sent by P i such that From PO(m 1 ), PO(m x ) and the send condition SC 1 (m dlv(m x ). Due to SC 1 (m Due to assumption bcast(m x mx Figure 4: Proof of Lemma 3.1. From the previous Lemma and from the definition of the width of a causal window given in Section 3 we have: Theorem 4.2 In a PAK causal broadcasting protocol based on SC 1 , i of CW j i is equal to 1 for any P i and P j . Proof i represents the number of consecutive messages m , sent by process P j (i.e., bcast(m x which have not arrived at process causally precede a message m just arrived at P i (i.e., bcast(m x would violate causal ordering with each message of the sequence . From the Lemma 4.1 if there is a causal ordering violation between m 0 and so the number of consecutive messages, sent by P j , that may violate causal ordering is at most one. Hence the claim follows. 2 Now, by considering the FIFO property of channels, we have the following Lemma: Lemma 4.3 Let b E be a distributed computation and m 2 M( b E) be a message sent by process P i . There is at most one message m x 2 M( b E) concurrent to m for each process P j (i 6= j). Proof If m and m x are concurrent follows that bcast(m x From PO(m x ), PO(m) and the channel FIFO property, the following sequence of events occurs in P j and P i as shown in Figure 5.a: Suppose there is another message m x 0 sent by P j and concurrent to m. From the send condition and from the definition of concurrent messages given above we have: bcast(m x 0 (a) (b) ackm ackm mx ackmx ackmx Figure 5: Proof of Lemma 3.3. shown in Figure 5.b. From FIFO property and PO(m x 0 ), on process process P j we have: dlv(m) OE j a:ack(m x 0 ), it follows that: bcast(m x which contradicts the send condition SC 1 (m x at most one message, sent by each distinct process, can be concurrent to m. 2 From the previous Lemma and the definition of the width of a causal window given in Section 3 we have: Theorem 4.4 In a PAK causal broadcasting protocol based on SC 1 , CO j i is equal to 1 for any P i and P j . Proof CO 2 i represents the number of consecutive messages m , sent by process P j which have been delivered to P i and are concurrent to a message m just arrived at P i . From Lemma 4.3, there is at most one message m x concurrent to m for each process P j (i 6= j), so is equal to one and the claim follows. 2 Hence, the range of variation of all the causal windows is limited and the step by which each vector time counter increases is 1 (due to the FIFO property of channels). So we have the following invariant: This allows a modulo k implementation of counters of vector times by choosing k greater than the maximum difference between any two counters i.e., k - 3. An example of such a window is shown in Figure 6 with k equal to 4. A remark on sliding windows. From Figure 6 it can been devised that a causal window is a particular type of sliding window. The sliding window is a technique widely used for flow-control in point-to-point data transfer protocols to avoid loss of messages and to ensure FIFO deliveries message delivered message delayed Figure over asynchronous and unreliable communication systems (e.g., TCP) [11, 17, 23]. This technique induces a closed loop between sender and receiver which allows not to overload buffer spaces of the receiver and to avoid network congestion by controlling the transmission rate of the sender. So the interest of a causal window lies also in the fact that it could be used as a part of a flow-control layer of a group communication system to provide causal communication to the above layer. 4.3 The Protocol The behavior of a process P i when executing the PAK protocol is described in Figure 7. When requesting to broadcast a message m, process P i first waits till a previous broadcast message, if any, is fully acknowledged (line S1) and then sets the variable processing broadcast to TRUE, stores the current vector time V T i in V Tm and sends m with attached V Tm as an atomic action (lines S2-S3). Afterwards, it waits till message m be fully acknowledged (i.e., an ack m message arrives from each member of P ) and, then, the local timestamp V T i [i] is increased by one module k (line S5) and finally successive broadcast messages are enabled by resetting processing broadcast (line S6). Upon arrival at process P i of message m, its delivery is determined by its delivery condition ng In this particular case, CP h i is equal to one. Message m is delivered as soon as the predicate DC(m) is true (line R1). When a broadcast message m, sent by P j , is delivered (line R2) to P i , the vector time V T i [j] is updated (line R3) and an ack m message is sent to P j (line R4). Lines up to R4 are executed atomically. 7 Note that (\Gammah) mod and that by definition of causal window k ? h. init processing broadcast := FALSE; for each h 2 procedure BCAST(m; is the message, P i is the sender % begin wait (:processing broadcast); (S1) processing broadcast := TRUE; V Tm for each h 2 ng do send (m; V Tm) to P h ; od % event bcast(m) % (S3) wait (for each h 2 ng do (ack m ) arrives from P h ; od); (S4) processing broadcast := FALSE; % event f:ack(m) % (S6) end. when (m;V Tm) arrives at P i from event arr(m) % begin ng event s:ack(m) % (R4) end. Figure 7: The PAK broadcasting protocol based on causal windows 4.4 Correctness Proof Theorem 4.5 Delivery events respect causal ordering (Safety). Proof Let us consider two messages m 1 and m 2 sent by processes P i and P j respectively and delivered to P h out of causal order (i.e., bcast(m 1 Lemma 4.1 we have that line of the protocol, we get: V Tm2 Upon the delivery of m 2 to P h , the delivery condition DC(m 2 ) (line R1) requires one of the following conditions be true: 1. [i], that is V Tm 1 2. V Tm2 k, that is V Tm1 By definition and line R2, V T h [i] contains the number of messages sent by P i and delivered to P h . Successive messages, sent by P i , are delivered in FIFO order by the assumption on FIFO channels and by the send condition SC 1 . As m 1 is delivered (by hypothesis) we have either V Tm1 mod k. In both cases, by considering conditions 1 and 2, m 1 has already been delivered; this contradicts the hypothesis that m 1 was delivered after m 2 . 2 Theorem 4.6 Each message will be eventually delivered (Liveness). Proof Let m x be the x-th message sent by P i and arrived at P j but never delivered. Given the delivery condition DC(m x ) of line R1, it follows: Two cases have to be considered: From the send condition SC(m upon the arrival of message m x at process P j , all messages m sent by P i were delivered to P j . So by line R2, after the delivery of m is equal to x \Gamma 1 which contradicts (P1). There must be at least one message m, causally preceding m x , sent by P k that either has not arrived at process P j or is arrived and delayed, so the delivery of m x would violate the causal ordering. From Lemma 4.1, we have Considering the reliable and broadcast nature of channels and messages respectively, sooner or later m arrives at P j . Now, two cases are possible: 8 This part of the proof is similar to the one in [9]. 1. m is delivered. This causes, by line R2, the delivery of m x . 2. m is delayed. The same argument can be applied to a message m x 0 , sent by Pw (with h), such that m x 0 ;m. Due to the finite number of processes and messages, sooner or later we fall either in case 1 or in the case we have a contradiction.5 Variants of the Protocol This section shows two variants of the previous protocol. The first one allows a process to have a certain number of outstanding unacknowledged messages at any time (credit). The second assumes that only a subset of messages be acknowledged (positive/negative acknowledgement approach). The aim of both variants is to reduce the number of local synchronizations of the protocol due to the send condition. 5.1 A PAK Broadcast Protocol using Credits Here we suppose that processes have a credit of ct - 1, i.e., a process P i can send up to ct - 1 consecutive broadcast messages m before receiving the corresponding acks. So, the send condition SC 1 can be extended as follows: ct (m Credits potentially reduce the number of synchronization in the send condition, but increase the width of the causal window 9 . Indeed, as shown in Figure 8.a, upon the arrival of a message m sent by P j , there could be at most ct consecutive messages sent by P k that causally precede m (bcast(m x ct \Gamma 1g)) and such that dlv(m) OE i dlv(m x ) (with ct \Gamma 1g). Let us assume m ct be a message sent by P k that causally precedes message m and that dlv(m) OE i dlv(m ct ). It follows dlv(m ct ), bcast(m ct and, because of, dlv(m) OE i dlv(m 0 ), we have bcast(m ct ct ), an absurdity. So upon the arrival of message m at process P i , we have V Tm ct. In figure 8.a, for clarity's sake, only the ack messages that produce the fully:ack events are depicted. 9 The local synchronization SC ct only weakens SC1 . In fact the number of local synchronizations is the same. How- ever, if the credit is appropriately chosen (as a function of the network latency), broadcast of very few messages will be prohibited because acks will be received before the credit is exhausted. So the number of "real" synchronizations is actually reduced. ct ct (a) (b) ct ct Figure 8: An example of message scheduling of a protocol with credits. On the other hand, as shown in Figure 8.b, upon the arrival of a message m sent by P j , there can be at most ct consecutive messages sent by P k and such that bcast(m y (with y ct \Gamma 1g). Indeed, due to the FIFO property, message m will be delivered to P k before the arrival at P j of the ack message related to m 0 . So bcast(m) ! bcast(m ct ). Hence, upon the arrival of message m at process P i , we have Hence, concerning the width of the causal windows, we have the following invariant: So a modulo k implementation of vector times with k - 2ct allowed and the size of such vectors is ndlog bits. To manage credits the protocol of Figure 7 needs some modification. In particular, to implement the send condition SC ct , the boolean variable processing broadcast becomes an integer one (initialized to zero) and lines S1 and S2 should be replaced with the following lines: wait (processing broadcast ! ct); processing broadcast := processing broadcast and lines S6 becomes: processing broadcast := processing broadcast \Gamma Finally, for the delivery condition, line R1 will be replaced ng ct ct Figure 9: Impossibility of a causal violation with more than ct consecutive messages. A remark on memory requirements. Up to now we have considered the causal window width as a function of the credit of the sender. Thus, the protocol delays all messages arrived too early at a process (the maximum number of pending messages is ct(n \Gamma 1)). If the buffer of the receiver has not enough space, it overflows dropping incoming messages. This situation can be mastered by associating a credit with the receiver. Let us define for a process P i a width of a causal window ct (being wt the credit of the receiver) such that 1 - wr - ct. An arriving message m, whose V Tm [j] fall in the interval [V T will be stored, delayed, delivered and acknowledged by P i . If its V Tm [j] falls in the interval [V T will be discarded by P i without sending the ack message. Managing a credit associated with the receiver requires, then, that each receiver has mechanisms to remove message duplication and each sender has a timer which triggers retransmission of messages if no ack is received within a deadline. A discussion about the use of previous mechanisms to support causal windows can be found in [3]. 5.2 A PAK/NAK Broadcast Protocol The protocol of Section 4.3 can be easily adapted for a solution using a PAK/NAK acknowledgment reducing the number of synchronization among messages due to the send condition compared to a PAK one. We assume that a process P i can send a sequence of ct messages without being acknowledged and then an ack message is required from the other processes after the receipt of the ct-th message sent by P i . A process is allowed to send the ct 1-st message only when the ct-st message has been fully acknowledged. In this case, the send condition becomes: if ct As for the protocol using credits, the width of the causal windows is 2ct. A violation of causal ordering will be always in the current range of the causal window since a message m sent by P i can create a causality violation in process P k at most with ct consecutive messages sent by P j . The acks required by the (i mod ct ct \Gamma 1)-st messages avoid a causality violation including more than ct consecutive messages. In particular, Figure 9 shows that if a message m, sent by P j , creates a causality violation in process P k with ct consecutive messages sent by P i then an absurdity follows, (depicted by thick arrows). Using the same argument of Section 5.1, no more of ct messages can be concurrent to any message of the computation due to FIFO property of channels. To implement a PAK/NAK protocol, the delivery condition is the same as the one of the protocol of Section 5.1 and the protocol of Figure 7 needs the following modifications. Lines S1 and S2 should be replaced with the following ones: (:processing broadcast); ct processing broadcast := TRUE; and line S6 becomes: processing broadcast := FALSE; 6 Conclusion In this paper a PAK causal broadcasting protocol based on causal windows have been proposed. A causal window actually represents the range of variation of vector time counters in the delivery condition of a causal ordering protocol. This protocol allow a modulo k implementation of vector times when considering k greater than the width of any causal window. This has been achieved by exploiting the causal information implicitly carried by ack messages. Compared with protocols that does not use control messages [5, 9, 18], the cost we pay is a little computational overhead and the presence of local synchronizations between messages, sent by the same process, due to the send condition (which reduces the potential concurrency of the protocol). To reduce the number of local synchronizations, we have discussed two variations of the pro- tocol. In both variations the reduction of the number of local synchronizations is payed by wider causal windows. The first variation allows a process to transmit a certain number of successive messages before receiving the corresponding acknowledgements (credit). This solution only potentially reduces the number of local synchronizations. However, if the credit is appropriately chosen, broadcast of very few messages will be prohibited because acks will be received before the credit is exhausted. At the same time, if the credit is not too high, the difference between sending n integer as a vector clok and ndlog ke bits as vector clock will be significant. So a proper choice of a credit value will lead to overhead reduction and insignificant loss of concurrency. The second variation employs a positive/negative method, i.e., it requires a local synchronizations between a message m x and the succesive message, sent by the same process, only if x is a multiple of a predefined parameter. This solution reduces the number of local synchronization and the message traffic generated by the protocol. Compared to the first variation, this seems to be well suited for high latencies networks. In the paper we also showed how the notion of causal window is related to the one of sliding windows used for FIFO flow control, a local synchronization due to the send condition is strictly connected to the concept of stability of a message in a group of processes and how this protocol can be adapted to avoid buffer overflow. The interested reader can find a causal broadcasting protocol based on causal window well suited for unreliable network in [3]. The description includes additional data structures and mechanisms, a process has to endow, in order to avoid lost of messages and message duplications. Acknowledgments The author would like to thank Ken Birman, Bruno Ciciani, Roy Friedman, Achour Mostefaoui, Michel Raynal, Ravi Prakash, Mukesh Singhal and Robbert Van Renesse for comments and many useful conversations on the work described herein. The author also thanks the anonymous referees for their detailed comments and suggestions that improved the content of the paper. --R "Causal Memory: Definitions, Implementation, and Programming" "Transis; a Communication Subsystem for High Availability" "A Positive Acknowledgement Protocol for Causal Broadcasting" "The Hierarchical Daisy Architecture for Causal Delivery" "Causal Deliveries of Messages with Real-Time Data in Unreliable Networks" "Efficient \Delta-causal Broadcasting" "A Note on Reliable Full-Duplex Transmission Over Half-Duplex Links" "Reliable Communication in the Presence of Failures" "Lightweight Causal Order and Atomic Group Multicast" "Reliable Distributed Computing with the ISIS Toolkit" "A Protocol for Packet Network Interconnection" "Logical Time in Distributed Computing Systems" "A Survey of Light-Weight Protocols for High-Speed Networks" "Time, Clocks and the Ordering of Events in a Distributed System" "Virtual Time and Global States of Distributed Systems" "AnAdaptive Causal Ordering Algorithm Suited to Mobile Computing Environments" "Networks and Distributed Computation" "The Causal Ordering Abstraction and a Simple Way to Implement It" "Securing Causal Relationship in Distributed Systems" "Logically Instantaneous Message-Passing in Asynchronous Distributed Systems" "A New Algorithm Implementing Causal Ordering" "A Data Transfer Protocol" "Computer Networks" --TR --CTR Roberto Baldoni, Response to Comment on "A Positive Acknowledgment Protocol for Causal Broadcasting", IEEE Transactions on Computers, v.53 n.10, p.1358, October 2004 Giuseppe Anastasi , Alberto Bartoli , Giacomo Giannini, On Causal Broadcasting with Positive Acknowledgments and Bounded-Length Counters, IEEE Transactions on Computers, v.53 n.10, p.1355-1358, October 2004

sliding windows;causal broadcasting;happened-before relation;distributed systems;group communication;asynchrony;vector times

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Algebraic and Geometric Tools to Compute Projective and Permutation Invariants.

AbstractThis paper studies the computation of projective invariants in pairs of images from uncalibrated cameras and presents a detailed study of the projective and permutation invariants for configurations of points and/or lines. Two basic computational approaches are given, one algebraic and one geometric. In each case, invariants are computed in projective space or directly from image measurements. Finally, we develop combinations of those projective invariants which are insensitive to permutations of the geometric primitives of each of the basic configurations.

Introduction Various visual or visually-guided robotics tasks may be carried out using only a projective representation which show the importance of projective informations at different steps in the perception-action cycle. We can mention here the obstacle detection and avoidance [14], goal position prediction for visual servoing or 3D object tracking [13]. More recently it has been shown, both theoretically and experimentally, that under certain conditions an image sequence taken with an uncalibrated camera can provide 3-D Euclidean structure as well. The latter paradigm consists in recovering projective structure first and then upgrading it into Euclidean structure [16,4]. Additionally, we believe that computing structure without explicit camera calibration is more robust than using calibration because we need not make any (possibly incorrect) assumptions about the Euclidean geometry (remembering that calibration is itself often erroneous). All these show the importance of the projective geometry as well in computer vision than in robotics, and the various applications show that projective informations can be useful at different steps in the perception-action cycle. Still the study of every geometry is based on the study of properties which are invariant under the corresponding group of transformations, the projective geometry is characterized by the projective invariants. This paper is dedicated to the study of the various configurations of points and/or lines in 3D space; it gives algebraic and geometric methods to compute projective invariants in the space and/or directly from image measurements. In the 3D case, we will suppose that we have an arbitrary three-dimensional projective representation of the object obtained by explicit projective reconstruction. In the image case, we will suppose that the only information we have is the image measurements and the epipolar geometry of the views and we will compute three-dimensional invariants without any explicit reconstruction. First we show that arbitrary configurations of points and/or lines can be decomposed into minimal sub-configurations with invariants, and these invariants characterize the the original configuration. This means that it is sufficient to study only these minimal configurations (six points, four points and a line, three points and two lines, and four lines). For each configuration we show how to compute invariants in 3D projective space and in the images, using both algebraic and geometric approaches. As these invariants generally depend on the order of the points and lines, we will also look for features which are both projective and permutation invariants. Projective Invariants Definition 1. Suppose that p is a vector of parameters characterizing a geometric configuration and T is a group of linear transformations, such that ae are y are vectors of homogeneous coordinates. A function I(p) is invariant under the action of the group T if is the value of I after the transformation T. If I 1 are n invariants, any f(p) = f(I variant. So if we have several invariants, it is possible that not all of them are functionally independent. The maximum number of independent invariants for a configuration is given by the following proposition [5,6]: Proposition 2. If S is the space parameterizing a given geometric configuration (for example six points, four lines) and T is a group of linear transformations, the number of functionally independent invariants of a configuration p of S under the transformations of T is: where Tp is the isotropy sub-group of the configuration p defined as aepg. Generally, minp2S (dim(Tp certain types of configurations have non trivial isotropy sub-groups. For example, for Euclidean transformations, the distance between two 3D points is an invariant and dim(S) \Gamma dim(T Consequently, minp2S (dim(Tp ))) 6= 0. Indeed, the sub-group of rotations about the axis defined by two points, leaves both of the points fixed. The most commonly studied transformation groups in computer vision are the Euclidean, affine and projective transformations. Since we want to work with weakly calibrated cameras, we will use projective transformations. As projective transformations of 3D space have points 3 and lines 4, using proposition 2, we can easily see that we need at least six points, four points and a line, two points and three lines, three points and two lines or four lines to produce some invariants. We will say that these configurations are minimal. Taking a non-minimal configuration of points and lines, we can decompose it into several minimal configurations. It is easy to show that the invariants of the sub configurations characterize the invariants of the original configuration. This means that we only need to be able to compute invariants for minimal configu- rations. For example, consider a configuration of seven points denoted M i;i=1::7 . From proposition 2 there are 3 \Theta 7 \Gamma independent invariants. To obtain a set of six independent invariants characterizing the configuration it is sufficient to compute the three independent invariants - i;i=1::3 of the configuration M i;i=1::6 and the three independent invariants - i;i=1::3 of the configuration M i;i=2::7 . We only discuss configurations of points and lines, not planes. Invariants of configurations of planes and/or lines can be computed in the same way as those of points and/or lines, by working in the dual space [12,2]. For the same reasons, we do not need to consider in detail configurations of two points and three lines. Indeed, these configurations define six planes which correspond to configurations of six points in the dual space [2]. The other four minimal configurations we will study in detail are: six points, four points and a line , three points and two lines and four lines. 1.1 Projective Invariants Using Algebraic Approach Consider eight points in space represented by their homogeneous coordinates compute the following ratio of determinants: k denotes the value oe(k) for an arbitrary permutation oe of f1; The invariant I can be computed also from a pair of images using only image measurements and the fundamental matrix using the Grassmann-Cayley, also called the double algebra as below [1,3]: and (fi ? stands for the expression sign(fi ? Note that if we change the bases in the images such that ff are the image coordinates in the new bases, the corresponding fundamental matrix is F so the quantities are independent of the bases chosen. 1.2 Six Point Configurations Now consider a configuration of six points A i;i=1::6 in 3D projective space. From proposition 2, there are 3 independent invariants for this configuration. Using (2) and (3) we can deduce the following three invariants: I 13\Gamma26 34\Gamma25 I 14\Gamma36 I To show that they are independent, change coordinates so that A i;i=1::5 become a standard basis. Denote the homogeneous coordinates of A 6 in this basis by Computing I j;j=1::3 in this basis we obtain I s and I s . These invariants are clearly independent. Alternatively, one can also take a geometric approach to compute six point invariants. The basic idea is to construct cross ratios using the geometry of the configuration. In this case we will give two different methods that compute independent projective invariants. The first method constructs a pencil of planes using the six points. Taking two of the points, for example A 1 and A 2 , the four planes defined by A 1 ; A 2 and A k;k=3::6 belong to the pencil of planes through the line A 1 A 2 . Their cross ratio is an invariant of the configuration. Taking other pairs of points for the axis of the pencil gives further cross ratios. The relation between these and I j;j=1::3 is: The second method [5] consists in constructing six coplanar points from the six general ones and computing invariants in the projective plane. For exam- ple, if take the plane A 1 A 2 A 3 and cut it by the three lines A 4 A 5 , A 4 A 6 and A 5 A 6 obtaining the intersections M 1 , M 2 and M 3 coplanar with A i;i=1:3 . Five coplanar points, for example A i;i=1::3 ; M i;i=1;2 , give two cross ratios Any other set of five, for example A i;i=1::3 ; M i=1;3 , gives two further cross only three of the four cross ratios are independent. Indeed The relation between - i;i=1::3 and I j;j=1::3 are - I3 and - So we have several methods of computing geometric invariants in 3D space. To do this, we need an arbitrary projective reconstruction of the points. However, the invariants can be also be computed directly from the images by using the fact that the cross ratio is preserved under perspective projections. First consider the case of the pencil of planes. We know that the cross ratio of four planes of a pencil is equal to the cross ratio of their four points of intersections with an arbitrary transversal line. So if we are able to find the image of the intersection point of a line and a plane we can also compute the required cross ratio in the image. The coplanar point method uses the same principle, computing the intersection of a line and a plane from image measurements. We want to compute the images of the intersection of a line A 3 A 4 and a plane only the projections of the five points in two images a i;i=1::5 and a 0 i;i=1::5 and the fundamental matrix between the two images. Take a point p in the first image and a point p 0 in the second one. These points are the projections we are looking for if and only if 3 \Theta a 0 verify (details in [2]): (a 0 2 \Theta a 0 1 \Theta (Fp \Theta (a 0 3 \Theta a 0 F (a 2 \Theta a 5 ) \Theta (a 1 \Theta p) The first equation is linear. The second one is quadratic, but it is shown in [2] that it can be decomposed into two linear components, one of which is irrelevant (zero only when p belongs the epipolar line of a 0 finally we obtain two linear equations which give the solution for p and 3 \Theta a 0 Another way to compute the intersection of a plane and a line from image measurements is to use the homography induced between the images by the plane. Let us denote the homography induced by the 3D plane \Pi by H. The image n 0 of the intersection N of a line L and the plane \Pi is then To compute the homography H of the plane A 1 A 2 A 5 , we use the fact that ae i a 0 We denote a a 0 and the coefficients of the matrix H by h ij . Then for each j 2 f1; 2; 5; eg, we have: As H is defined only up to a scale factor, it has only eight independent degrees of freedom and can be computed from the eight equations of (5). 1.3 Configurations of One Line and Four Points Denoting the four points by A i;i=1::4 and the line by L we obtain (cf. (2) and (3)) the following invariant: arbitrary two distinct points on the line L, ff are the projections of A i;i=1::4 and L in the two images, Using the geometric approach, the four planes LA i;i=1::4 belong to a pencil so they define an invariant cross ratio -. Another approach, given by Gros in [5] is to consider the four planes defined by the four possible triplets of points and cut the line L with them. This gives another cross ratio - 0 for the configuration. Of course, we only have one independent invariant, so there are relations between I , - and - 0 . Indeed, we have The method of computing - and - 0 directly from the images is basically the same as for the configuration of six points . 1.4 Configurations of Three Points and Two Lines The following two ratios are independent invariants for configurations of three points A i;i=1::3 and two lines L k;k=1;2 in 3D: I I These invariants can also be obtained as follows. Cut the two lines with the plane defined by the three points to give R 1 and R 2 . The five coplanar points define a pair of invariants, for example Another way to compute a pair of independent invariants for this configuration is to consider the three planes L 1 A i;i=1::3 and the plane A 1 A 2 A 3 and cut them by L 2 This gives a cross ratio - 0 -2 . Changing the role of L 1 and L 2 gives another cross ratio - 0 . The cross ratios - i;i=1;2 and - 0 i;i=1;2 can be computed directly in the images in the same way as the cross ratios of the configuration of six points (finding images of intersections of lines and planes). Configurations of Four Lines Consider four lines L i;i=1::4 in 3D projective space. This configuration has rameters, so naively we might expect to have independent invariant. However, the configuration has a 1D isotropy subgroup, so there are actually two independent invariants [6,5,2]. The existence of two independent cross ratios can also be shown geometri- cally. Assume first that the lines are in general position, in the sense that they are skew and none of them can be expressed as a linear combination of three others. Consider the first three lines L i;i=1::3 . As no two of them are coplanar, there is a one parameter family K of lines meeting all three of them. K sweeps out a quadric surface in space, ruled by the members of K and also by a complementary family of generators L, to which L 1 , L 2 and L 3 belong [11,6]. Members of K are mutually skew, and similarly for L, but each member of K intersect each member of L exactly once. Another property of generators is that all members of each family can be expressed as a linear combination of any three of them. By our independence assumptions, the fourth line L 4 does not belong to either family (if L 4 belonged to L it would be a linear combination of the L i;i=1::3 and if L 4 was in K it would cut each of the lines L i;i=1::3 ). Hence, L 4 cuts the surface in two real or imaginary points, A 4 and B 4 (these may be identical if L 4 is tangent to the surface). For each point of the surface there is a unique line of each family passing through it. Denote the lines of K passing through A 4 and respectively. Let these lines cut the L i;i=1::3 in A i;i=1::3 and respectively. In this way, we obtain two cross ratios: Before continuing, consider the various degenerate cases. - Provided that the first three lines are mutually skew, they still define the two families K and L. The fourth line can then be degenerate in the following ways: ffl If L 4 intersects one (L i ) or two (L of the three lines, we have the same solution as in the general case with A respectively. ffl If L 4 2 K intersects all three lines L i;i=1::3 , the two transversals T 1 and are equal to L 4 and the cross ratios are no longer defined. ffl If L 4 2 L (linear combination of the lines L i;i=1::3 ) every line belonging to K cuts the four lines L i;i=1::4 and there is only one characteristic cross ratio. Given two pairs of coplanar lines, say L there are two transversals intersecting all four lines. One is the line between the intersection of L 1 and L 2 and the intersection of L 3 and L 4 . In this case the cross ratio 1. The other is the line of intersection of the planes defined by defining a second cross ratio [5]. - If three of the lines are coplanar, say L i;i=1::3 , the plane \Pi containing them intersects L 4 in a point A 4 . If the lines L i;i=1::3 belong to a pencil with center the lines L i;i=1::3 and BA 4 define one cross ratio. Otherwise, we have no invariants. Finally, if all of the lines lie in the same plane \Pi, the possible cases are that they belong to a pencil and define one cross ratio or they do not belong to a pencil and there are no invariants. As in the preceding cases, we can also compute the invariants using the algebraic I I where l i and l 0 i are the projections of the L i;i=1::4 in the images, fi l i . The relation between I i;i=1;2 and - j;j=1;2 is I and I conversely, given I 1 and I 2 , - 1 and - 2 are the solutions of the equation As I 1 and I 2 are real, - 1 and - 2 are either real or a complex conjugate pair. Computation of - 1 and - 2 in 3D: First, we suppose that we know a projective representation of the lines in some projective basis. They can be represented by their Pl-ucker coordinates L We look for lines T that intersect the L i;i=1::4 . If it represents a transversal line if and only if: ae l (i) This is a homogeneous system in the six unknowns p i;i=1::6 , containing four linear equations and one quadratic one. It always has two solutions: (p k k=1;2 defined up to a scale factor, which may be real or complex conjugates. These lines cut the L i;i=1::4 in A i;i=1::3 and B i;i=1::3 which Computation of - 1 and - 2 from a Pair of Views: Define F to be the fundamental matrix and l i ) to be the projections of the lines L i;i=1::4 into the images. Consider a line l in the first image, with intersections x l with the lines l i;i=1::4 . If the intersections i of the epipolar lines with the l 0 i;i=1::4 in the second image all lie on some line l 0 , then l and l 0 are the images of a transversal line L. The condition of this can be written: ae This system of equations is difficult to deal with directly. To simplify the computation we make a change of projective basis, so that the new basis satisfies The fundamental matrix then becomes and the two pencils of epipolar lines take the form Using this parameterization, and after simplifications [2] we obtain: complex conjugates, are the significant 1 solutions of the system: ae with Projective and Permutation Invariants The invariants of the previous sections depend on the ordering of the underlying point and line primitives. However, if we want to compare two configurations, we often do not know the correct relative ordering of the primitives. If we work with order dependent invariants, we must compare all the possible permutations of the elements. For example, if I k;k=1::3 are the three invariants of a set of points A i;i=1::6 and we would like to know if another set B i;i=1::6 is projectively equivalent, we must compute the triplet of invariants for each of the 6!=720 1 The other two solutions are (1; 0; our parameterization we can consider without lost of generality that t i possible permutations of the B i . This is a very costly operation, which can be avoided if we look for permutation-invariant invariants. 3 Definition 3. A function I is a projective and permutation invariant for a configuration formed by elements only if we have for all projective transformations T and for all permutations - 2 S k The projective invariants we computed were all expressed as cross ratios, which are not permutation invariants. As the group of permutations S 4 has 24 elements, there are potentially 24 possible cross ratios of four collinear points. However, not all of these are distinct, because the permutations - have the same effect on the cross ratio. We know that all permutations can be written as products of transpositions of adjoining elements. For S 4 there are three such transpositions 4). The effect of - 1 and - 3 is F 1 - and that of - 2 is F 2 -. If we compute all the possible combinations of these functions, we obtain the following six values: - . To obtain permutation invariant cross ratios of four elements, it is sufficient to take an arbitrary symmetric function of the - i;i=1::6 . The simplest two symmetric functions J are not interesting because they give constant values. Taking the second order basic symmetric functions unbounded functions. Further invariants can be generated by taking combinations of the former ones proposed in [8] bounded between 2 and 2.8 or J2 bounded between 2 and 14. These functions are characterized by the fact that they have the same value for each of the six arguments - i;i=1::6 . Let us see what happens in the case of six points in projective space. We saw that the invariants of the points A i;i=1::6 given by (4) correspond to cross ratios of pencils of planes. Define: I A -(1) A -(2) A g. It is easy to see that I I 231456 , I I 132456 and I We are interested in the effect of permutations of points on the invariant . The first remark is that if we interchange the first two elements the value is the same: I which is obvious from a geometric viewpoint as A 1 A 2 A k and A 2 A 1 A k represent the same plane. If we fix the first two points and permute the last four, we permute the four planes giving the cross ratio. So if we apply one of the above symmetric functions, for example J , the results will be invariant. Using the following proposition proved in [2] we can find all the possible values for J(I -(1)-(2)-(3)-(4)-(5)-(6) ). 3 We generalize here the work of Meer et al. [8] and Morin [9] on the invariants of five coplanar points to 3D configurations of points and lines. Proposition 4. An arbitrary permutation - of S 6 can be written as a product of four permutations - 1k - 2n - 1m , is a permutation of the last four elements. Denote the value of I -(1)-(2)-(3)-(4)-(5)-(6) by -(I 123456 ). We have then -(I 123456 applying (for example) J , we obtain: because - changes only the order of the four planes giving the cross ratio. But - 1k has no effect on the value of the cross ratio: - 11 is the identity and - 12 interchanges the first two elements but leaves the planes invariant, so J(-(I 123456 J((- 2n (- 1m (I 123456 ))). Consequently, we obtain only C 2 different possible values instead of These will be denoted by Jmn = J((- 2n (- 1m (I 123456 ))). It is easy to see that ng. The geometric meaning of this is that we fix the pair of points giving the axis of the pencil of planes. As a function of I 1 ; I 2 and I 3 , the 15 values are: Note that the values of I 1 ; I 2 ; I 3 depend on the order of the points, but any other ordering gives the same values in different ordering. For this reason, we sort the 15 values after computing them to give a vector that characterizes the configuration independently of the order of underlying points. Now, consider the configuration of a line and four unordered points. This case is a very simple case because we have a cross ratio, so it is enough to take the value of the function J(I) where I is the invariant given by (6). Similarly to compute permutation invariants of four lines in space, we apply J to the pair of cross ratios - 1 and - 2 given in the section 1.5. It turn out that the order of the two sets of intersection points of the four lines with the two transversals change in the same way under permutations of the lines. When - 1 and - 2 are complex conjugates, J(- 1 ) and J(- 2 ) are also complex conjugates, but if we want to work with real invariants it is sufficient to take J The fourth basic configuration is the set of three points and two lines. We saw that we can compute invariants for this configuration using the invariants of five coplanar points. In this case it is sufficient to apply the results of [8], which show that if I 1 and I 2 are a pair of invariants for five coplanar points we can take the sorted list with elements: J(I 1 ), to obtain a permutation invariant. But this means that we make mixed permutations of points and lines, which is unnecessary because when we want to compare two such configurations we have no trouble distinguishing lines from points. When we have five points in the plane A 1 A 2 A 3 we will require that the center points of the considered pencils of lines be chosen among A i;i=1::3 and not among the intersection points R j;j=1;2 (see section 1.4). In this way we have the cross ratios of the form - . If we interchange A 1 and A 2 , we obtain I 2 . Interchanging A 2 and A 3 gives 1 I1 and interchanging A 1 and A 3 we obtain I1 I2 . The permutation of the two lines (i.e. R 1 and R 2 I1 . Hence, applying J we find that the sorted list of elements is a projective and permutation invariant for this configuration. 2.1 Some Ideas Concerning More Complex Configurations Consider a configuration of N ? 6 points 4 From proposition 2, this configuration has 3N independent invariants. Denote the N points by A i;i=1::N and the 3(N \Gamma 5) invariants of this configuration by - k 3 are the invariants of the sub-configurations A As the invariants - k i depend on the order of the points, we try to generalize the above approach to the N point case. First note that the invariants are cross ratios of four planes of a pencil I Ng. Consequently, for a given set fi; j; k; l; m;ng ae different values J(I ijklmn ). But, there are C 6 N subsets of six points, so we have C 6 6 different values which can be computed from the 3(N \Gamma 5) independent invariants - k We show how to obtain these values in the case of 7 points, but the approach is the same for N ? 7. Denote the independent invariants of this configuration by - 1 The C 6 values can be obtained by calculating for each subset of six points fi; j; k; l; m;ng ae 7g such that the values I I ikjlmn and I I ijklmn in function of - j subsets I I I 3 - 1 and then apply (10). Finally, we sort the resulting list of 105 elements. Conclusion We have presented a detailed study of the projective and permutation invariants of configurations of points, lines and planes. These invariants can be used as a basis for indexing, describing, modeling and recognizing polyhedral objects from perspective images. 4 We will consider only the case of points. The other cases can be handled similarly, but the approach is more complex and will be part of later work. The invariants of complex configurations can be computed from those of minimal configurations into which they can be decomposed. So it was sufficient to treat only the case of minimal configurations. Also, in projective space there is a duality between points and planes that preserves cross ratios, so configurations of planes or planes and lines can be reduced to point and point-line ones. For each configuration we gave several methods to compute the invariants. There are basically two approaches - algebraic and geometric - and in each case we showed the relations between the resulting invariants. We analyzed the computation of these invariants in 3D space assuming that the points and lines had already been reconstructed in an arbitrary projective basis, and we also gave methods to compute them directly from correspondences in a pair of images. In the second case the only information needed is the matched projections of the points and lines and the fundamental matrix. Finally, for each basic configuration we also gave permutation and projective invariants, and suggested ways to treat permutation invariance for more complicated configurations. Acknowledgment We would like to thank Bill Triggs for his carefully reading of the draft of this manuscript. --R Multiple image invariants using the double algebra. Modelisation projective des objets tridimensionnels en vision par ordi- nateur Computing three-dimensional projective invariants from a pair of images using the Grassmann-Cayley algebra From projective to euclidean reconstruction. 3D projective invariants from two images. Invariants of lines in space. Visually guided object grasping. Correspondence of coplanar features through p 2 Quelques contributions des invariants projectifs 'a la vision par ordina- teur. Th'ese de doctorat A comparison of projective reconstruction methods for pairs of views. Algebraic Projective Geometry. G'eom'etries affine Geometric invariants for verification in 3-d object tracking Applications of non-metric vision to some visual guided tasks A robust technique for matching two uncalibrated images through the recovery of the unknown epipolar geometry. Metric calibration of a stereo rig. --TR --CTR Franoise Dibos , Patrizio Frosini , Denis Pasquignon, The Use of Size Functions for Comparison of Shapes Through Differential Invariants, Journal of Mathematical Imaging and Vision, v.21 n.2, p.107-118, September 2004 Yihong Wu , Zhanyi Hu, Camera Calibration and Direct Reconstruction from Plane with Brackets, Journal of Mathematical Imaging and Vision, v.24 n.3, p.279-293, May 2006 Arnold W. M. Smeulders , Marcel Worring , Simone Santini , Amarnath Gupta , Ramesh Jain, Content-Based Image Retrieval at the End of the Early Years, IEEE Transactions on Pattern Analysis and Machine Intelligence, v.22 n.12, p.1349-1380, December 2000

indexation;uncalibrated stereo;grassmann-cayley algebra;projective reconstruction;projective and permutation invariants;cross ratio

298503

L-Printable Sets.

A language is L-printable if there is a logspace algorithm which, on input 1n, prints all members in the language of length n. Following the work of Allender and Rubinstein [SIAM J. Comput., 17 (1988), pp. 1193--1202] on P-printable sets, we present some simple properties of the L-printable sets. This definition of "L-printable" is robust and allows us to give alternate characterizations of the L-printable sets in terms of tally sets and Kolmogorov complexity. In addition, we show that a regular or context-free language is L-printable if and only if it is sparse, and we investigate the relationship between L-printable sets, L-rankable sets (i.e., sets A having a logspace algorithm that, on input x, outputs the number of elements of A that precede x in the standard lexicographic ordering of strings), and the sparse sets in L. We prove that under reasonable complexity-theoretic assumptions, these three classes of sets are all different. We also show that the class of sets of small generalized Kolmogorov space complexity is exactly the class of sets that are L-isomorphic to tally languages.

Introduction . What is an easy set? Typically, complexity theorists view easy sets as those with easy membership tests. An even stronger requirement might be that there is an easy algorithm to print all the elements of a given length. These "printable" sets are easy enough that we can efficiently retrieve all of the information we might need about them. Hartmanis and Yesha first defined P-printable sets in 1984 [HY84]. A set A is P-printable if there is a polynomial-time algorithm that on input 1 n outputs all of the elements of A of length n. Any P-printable set must lie in P and be sparse, i.e., the number of strings of each length is bounded by a fixed polynomial of that length. Allender and Rubinstein [AR88] give an in-depth analysis of the complexity of the P-printable sets. Once P-printability has been defined, it is natural to consider the analogous notion of logspace-printability. Since it is not known whether or not an obvious question to ask is: do the L-printable sets behave differently than the P-printable sets? In this paper, we are able to answer this question in the affirmative, at least under plausible complexity theoretic assumptions. Jenner and Kirsig [JK89] define L-printability as the logspace computable version of P-printability. Because L-printability implies P-printability, every L-printable set must be sparse and lie in L. In this paper we give the first in-depth analysis of the complexity of L-printable sets. (Jenner and Kirsig focused only one chapter on printability, and most of their printability results concern NL-printable sets.) y Department of Computer Science, University of Chicago, Chicago, IL 60637. The work of this author was supported in part by NSF grant CCR-9253582. z Department of Computer Science, University of Kentucky, Lexington, KY 40506-0046. The work of these authors was supported in part by NSF grant CCR-9315354. x DIMACS Center, Rutgers University, Piscataway, NJ 08855. The work of this author was supported by NSF cooperative agreement CCR-9119999 and a grant from the New Jersey Commission on Science and Technology. - The work of this author was supported in part by a University of Kentucky Presidential Fellowship. L. FORTNOW, J. GOLDSMITH, M. LEVY, AND S. MAHANEY Whenever a new class of sets is analyzed, it is natural to wonder about the structure of those sets. Hence, we examine the regular and context-free L-printable sets. Using characterizations of the sparse regular and context-free languages, we show in x 4 that every sparse regular or context-free language is L-printable. (Although the regular sets are a special case of the context-free sets, we include the results for the regular languages because our characterization of the sparse regular languages is simple and intuitive.) We might expect many of the properties of P-printable sets to have logspace analogues, and, in fact, this is the case. In x 5 we show that L-printable sets (like their polynomial-time counterparts) are closely related to tally sets in L, and to sets in L with low generalized space-bounded Kolmogorov complexity. A set is said to have small generalized Kolmogorov complexity if all of its strings are highly compressible and easily restorable. Generalized time-bounded Kolmogorov complexity and generalized space-bounded Kolmogorov complexity are introduced in [Har83] and [Sip83]. Several researchers [Rub86, BB86, HH88] show that P-printable sets are exactly the sets in P with small generalized time-bounded Kolmogorov com- plexity. [AR88] show that a set has small generalized time-bounded Kolmogorov complexity if and only if it is P-isomorphic to a tally set. Using similar techniques, we show in x 5 that the L-printable sets are exactly the sets in L with small generalized space-bounded Kolmogorov complexity. We also prove that a set has small generalized space-bounded Kolmogorov complexity if and only if it is L-isomorphic to a tally set. In x 6, we note that sets that can be ranked in logspace (i.e., given a string x, a logspace algorithm can determine the number of elements in the set - x) seem different from the L-printable sets. For sparse sets, P-rankability is equivalent to P- printability. We show a somewhat surprising result in x 6, namely that the sparse L-rankable sets and the L-printable sets are the same if and only if there are no tally sets in P \Gamma L if and only if Are all sparse sets in L either L-printable or L-rankable? Allender and Rubinstein [AR88] show that every sparse set in P is P-printable if and only if there are no sparse sets in FewP \Gamma P. In x 6, we similarly show a stronger collapse: every sparse set in L is L-printable if and only if every sparse set in L is L-rankable if and only if there are no sparse sets in FewP \Gamma L if and only if Unlike L-printable sets, L-rankable sets may have exponential density. Blum (see [GS91]) shows that every set in P is P-rankable if and only if every #P function is computable in polynomial time. In x 6, we also show that every set in L is L-rankable if and only if every #P function is computable in logarithmic space. 2. Definitions. We assume a basic familiarity with Turing machines and Turing machine complexity. For more information on complexity theory, we suggest either [BDG88] or [Pap94]. We also assume a familiarity with regular languages and expressions and context-free languages as found in [Mar91]. We denote the characteristic function of A by -A . We use the standard lexicographic ordering on strings and let jwj be the length of the string w. (Recall that w - lex v iff jwj ! jvj or and, if i is position of the leftmost bit where w and v differ, w .) The alphabet all strings are elements of \Sigma . We denote the complement of A by A. The class P is deterministic polynomial time, and L is deterministic logarithmic space; remember that in calculating space complexity, the machine is assumed to have separate tapes for input, computation, and output. The space restriction applies only to the work tape. It is known that L ' P, but it is not known whether the two classes L-PRIN are equal. The class E is deterministic time 2 O(n) , and LinearSPACE is deterministic space O(n). Definition 2.1. A set A is in the class PP if there is a polynomial time nondeterministic Turing machine that, on input x, accepts with more than half its computations A. A function f is in #P if there is a polynomial time nondeterministic Turing machine M such that for all x, f(x) is the number of accepting computations of M(x). Allender[All86] defined the class FewP. FewE is defined analogously. Definition 2.2. [All86] A set A is in the class FewP if there is a polynomial time nondeterministic Turing machine M and a polynomial p such that on all inputs accepts x on at most p(jxj) paths. A set A is in the class FewE if there is an exponential time nondeterministic Turing machine M and a constant c such that on all inputs x, M accepts x on at most 2 cn paths. (Note that this is small compared to the double exponential number of paths of an exponential-time nondeterministic Turing machine.) Definition 2.3. A set S is sparse if there is some polynomial p(n) such that for all n, the number of strings in S of length n is bounded by p(n) (i.e., jS =n j - p(n) ). set T over alphabet \Sigma is a tally set if T ' foeg , for any character oe 2 \Sigma. The work here describes certain enumeration properties of sparse sets in L. There are two notions of enumeration that are considered: rankability and printability. Definition 2.4. If C is a complexity class, then a set A is C-printable if and only if there is a function computable in C that, on any input of length n, outputs all the strings of length n in A. Note that P-printable sets are necessarily in P, and are sparse, since all of the strings of length n must be printed in time polynomial in n. Since every logspace computable function is also computable in polynomial time, L-printable sets are also P-printable, and thus are also sparse. Definition 2.5. If C is a complexity class, then a set, A, is C-rankable if and only if there is a function r A computable in C such that r A (In other words, r A (x) gives the lexicographic rank of x in A.) The function r A is called the ranking function for A. Note that P-rankable sets are necessarily in P but are not necessarily sparse. Furthermore, a set is P-rankable if and only if its complement is P-rankable. Finally, note that any P-printable set is P-rankable. Definition 2.6. If C is a complexity class, then two sets, A and B, are C- isomorphic there are total functions f and g computable in C that are both one-one and onto, such that y, and f is a reduction from A to B, and g is a reduction from B to A. In order for two sets to be P-isomorphic, their density functions must be close to each other: if one set is sparse and the other is not, then any one-one reduction from the sparse set to the dense set must have super-polynomial growth rate. By the same argument, if one has a super-polynomial gap, the other must have a similar gap. A lexicographic (or order-preserving) isomorphism from A to B is, informally, a bijection that maps the ith element of A to the ith element of B and maps the ith element of A to the ith element of B. Note that in the definition of similar densities, the isomorphisms need not be computable in any particular complexity class. This merely provides the necessary condition on densities in order for the two sets to be P-isomorphic or L-isomorphic. Definition 2.7. Two sets, A and B, have similar densities if the lexicographic 4 L. FORTNOW, J. GOLDSMITH, M. LEVY, AND S. MAHANEY isomorphisms from A to B and from B to A are polynomial size bounded. The notion of printability, or of ranking on sparse sets, can be considered a form of compression. Another approach to compression is found in the study of Kolmogorov complexity; a string is said to have "low information content" if it it has low Kolmogorov complexity. We are interested in the space-bounded Kolmogorov complexity class defined by Hartmanis [Har83]. Definition 2.8. Let M v be a Turing machine, and let f and s be functions on the natural numbers. Then we define and M v uses s(n) space)g: Following the notation of [AR88], we refer to y as the compressed string, f(n) as the compression, and s(n) as the restoration space. Hartmanis [Har83] shows that there exists a universal machine M u such that for all v, there exists a constant c such that KS v [f(n); s(n)] ' KS u the subscript and let 3. Basic Results. We begin by formalizing some observations from the previous section. Observation 3.1. If A is L-printable, then A has polynomially bounded density, i.e., A is sparse. This follows immediately from the fact that logspace computable functions are P-time computable (i.e., L-printability implies P-printability), and from the observations on P-printable sets. Proposition 3.2 ([JK89]). If A is L-printable, then A 2 L. Proof. To decide x 2 A, simulate the L-printing function for A with input 1 jxj . As each y 2 A is "printed," compare it, bit by bit, with x. If accept. Because the comparisons can be done using O(1) space, and the L-printing function takes O(log jxj) space, this is a logspace procedure. Proposition 3.3. If A is L-rankable, then A 2 L. Proof. Note that the function x \Gamma 1 (the lexicographic predecessor of x) can be computed (though not written) in space logarithmic in jxj. Since logspace computable functions are closed under composition, r A can be computed in logspace, as can r A Proposition 3.4. If A is L-printable, then A is L-rankable. Proof. To compute the rank of x, we print the strings of A up to jxj and count the ones that are lexicographically smaller than x. Since A is sparse, by Observation 3.1, we can store this counter in logspace. We can now prove the following, first shown by [JK89] with a different proof. Proposition 3.5 ([JK89]). If A is L-printable, then A is L-printable in lexicographically increasing order. Proof. To prove this, we use a variation on selection sort. Suppose the logspace machine M L-prints A. Then we can construct another machine, N , to L-print A in lexicographically increasing order. Note that it is possible to store an instantaneous description of a logspace machine, i.e., the position of the input head, the state, the contents of the worktape, and the character just output, in O(log jxj) space. The basic idea is that we store, during the computation, enough information to produce three strings: the most recently printed string (in the lexicographically ordered printing), the current candidate for the next string to be printed, and the L-PRIN current contender. We can certainly store three IDs for M in logspace. Each ID describes the state of M immediately prior to printing the desired string. In addition to storing the IDs, we must simulate M on these three computations in parallel, so that we can compare the resulting strings bit by bit. If the contender string is greater than the last string output (so it has not already been output) and less than the candidate, it becomes the new candidate. Otherwise, the final ID of the computation becomes the new contender. These simulated computations do not produce output for N ; when the next string is found for N to print, its initial ID is available, and the simulation is repeated, with output. Using the same technique as in the previous proof, one can easily show the following Proposition 3.6. If A is L-printable, and A - =log B, then B is L-printable as well. 4. L-Printable Sets. We begin this section with a very simple example of a class of L-printable sets. Proposition 4.1 ([JK89]). The tally sets in L are L-printable. Proof. On input of length n, decide whether 1 n 2 A. If so, print it. One may ask, are all of the L-printable sets as trivial as Proposition 4.1? We demonstrate in the following sections that every regular language or context-free language that is sparse is also L-printable (see Theorem 4.5 and Corollary 4.14). We also give an L-printable set that is neither regular nor context-free (see Proposition 4.15). 4.1. Sparse Regular Languages. We show that the sparse regular languages are L-printable. In order to do so, we give some preliminary results about regular expressions. Definition 4.2 ([BEGO71]). Let r be a regular expression. We say r is unambiguous if every string has at most one derivation from r. Theorem 4.3 ([BEGO71]). For every regular language L, there exists an unambiguous regular expression r such that Proof. (Sketch) Represent L as the union of disjoint languages whose DFA's have a unique final state. Using the standard union construction of an NFA from a DFA, we get an NFA with the property that each string has a unique accepting path. Now, using state elimination to construct a regular expression from this NFA, the unique path for each string becomes a unique derivation from the regular expression. We should note that even though removal of ambiguity from a regular expression is, in general, PSPACE-complete [SH85], this does not concern us. Theorem 4.3 guarantees the existence of an unambiguous regular expression corresponding to every regular language, that is sufficient for our needs. We now define a restricted form of regular expression, that will generate precisely the sparse regular languages. (Note that a similar, although more involved, characterization was given in [SSYZ92]. They give characterizations for a variety of densities, whereas we are only concerned with sparse sets.) Definition 4.4. We define a static regular expression (SRE) on an alphabet \Sigma inductively, as follows: 1. The empty expression is an SRE, and defines ;, the empty set. 2. If x 2 \Sigma or string ), then x is an SRE. 3. If s and t are SREs, then st, the concatenation of s and t, is an SRE. 4. If s and t are SREs, then s + t, the union of s and t, is an SRE. 5. If s is an SRE, then s is an SRE iff: a) s does not contain a union of two SREs; and, 6 L. FORTNOW, J. GOLDSMITH, M. LEVY, AND S. MAHANEY b) s does not contain any use of the operator. Note the restriction of the operator in the above definition. I.e., can only be applied to a string. This is the only difference between SREs and standard regular expressions. We can alternately define an SRE as a regular expression that is the sum of terms, each of that is a concatenation of letters and starred strings. Theorem 4.5. Let R be an unambiguous regular expression. Then L(R) is sparse iff R is static. Proof. We first prove two lemmas about "forbidden" subexpressions. Lemma 4.6. Let ff, fi; S be non-empty regular expressions such that and S is unambiguous. Then there is a constant k ? 0 such that, for infinitely many k strings of length n. Proof. Let such that u 2 L(ff) and v 2 L(fi). Let S is unambiguous, there must be at least two strings of length k in L(S), namely u jvj and v juj . So, for any length n such that there are at least 2 strings of length n in L(S). Lemma 4.7. Let ff, fi; S be non-empty regular expressions such that S is unam- biguous, where S is either of the form (ff fi) or of the form (fffi ) . Then, there is a constant k such that, for infinitely many n, k strings of length n. Proof. Let such that u 2 L(ff) and v 2 L(fi). Suppose jvj. If S is unambiguous, there are at least two distinct strings of length k in L(S), namely, u jvj v and v juj+1 . So, for any length n such that there are at least 2 k strings of length n in L(S). The proof is very similar if unambiguous. It is clear that unambiguity is necessary for both lemmas. For example, the is not static, but L((a that is sparse. Note that if R is the empty expression, the theorem is true, since R is static, and which is certainly sparse. So, for the rest of the proof, we will assume that R is non-empty. To show one direction of Theorem 4.5, suppose R is not static. Then it contains a subexpression that is either of the form (fl or of the form (fl 0 ff In the first case, by a small modification to the proof of lemma 4.6, L(R) is not sparse. In the second case, by a similar modification to the proof of lemma 4.7, L(R) cannot be sparse. Now, suppose R is static. If contains only the string x. If is either a string of characters or a single character, L(R) can have at most one string of any length. are SREs. Let p r (n) and p s (n) bound the number of strings in L(r) and L(s), respectively. Then there are at most p r (n)+p s (n) strings of length n. Finally, suppose are SREs. Let p r (n) and p s (n) bound the number of strings in L(r) and L(s), respectively. Then, the number of strings of length n is: The degree of q is bounded by 1 on the complexity of R, L(R) is sparse. L-PRIN Note that the second half of the proof does not use unambiguity. Hence, any static regular expression generates a sparse regular language. Theorem 4.8. Let R be an SRE. Then L(R) is L-printable. Proof. Basically, we divide R into terms that are either starred expressions or non-starred expressions. For example, we would divide 0(1 into three parts: 0(1 internally L-print each term independently, and check to see if the strings generated have the correct length. In our example, to print strings of length 9, we might generate 0110, 11, and 0011, respectively, and check that the combined string is in fact 9 characters long. (In this case, the string is too long, and is not printed.) Let k be the number of stars that appear in R. Partition R into at most 2k subexpressions, k with stars, and the others containing no stars. The machine to L-print L(R) has two types of counters. For each starred subex- pression, the machine counts how many times that subexpression has been used. For a string of length n, no starred subexpression can be used more than n times. Each counter for a starred subexpression only needs to count up to n. Each non-starred subexpression generates only a constant number of strings. Thus, up to k +1 additional counters, each with a constant bound, are needed. (Note that the production may intermix the two types of counters, for instance if (x occurs.) The machine uses two passes for each potential string. First, the machine generates a current string, counting its length. If the string is the correct length, it regenerates the string and prints it out. Otherwise, it increments the set of counters, and continues. In this way, all strings of lengths - n are generated, and all strings of length n are printed. Lastly, we need to argue that this procedure can be done by a logspace machine. Each of the at most 2k must count up to n (for n sufficiently large, say, larger than jRj). Thus, the counting can be done in log n space. In addition, the actual production of a string requires an additional counter, to store a loop variable. The rest of the computation can be handled in O(1) space, using the states of the machine. Thus, L(R) is L-printable. Note that this L-printing algorithm may generate some strings in L(R) more than once. To get a non-redundant L-printer, simply modify the program to output the strings in lexicographic order, as in Proposition 3.5, or use an unambiguous SRE for L(R). Theorem 4.8 does not characterize the L-printable sets, as we see below. Proposition 4.9. There exists a set S such that S is L-printable and not regular Proof. The language L-printable (for any n, we print out only if n is even), but not regular. 4.2. Sparse Context-Free Languages. Using the theory of bounded context-free languages we can also show that every sparse context-free language is L-printable. Definition 4.10. A set A is bounded if there exist strings w such that Note the similarity between bounded languages and languages generated by SRE's. Note also that every bounded language is sparse. Ibarra and Ravikumar [IR86] prove the following. 8 L. FORTNOW, J. GOLDSMITH, M. LEVY, AND S. MAHANEY Theorem 4.11 ([IR86]). If A is a context-free language then A is sparse if and only if A is bounded. Ginsburg [Gin66, p. 158] gives the following characterization of bounded context-free languages. Theorem 4.12 ([Gin66]). The class of bounded context-free languages is the smallest class consisting of the finite sets and fulfilling the following properties. 1. If A and B are bounded context-free languages then A [ B is also a bounded context-free language. 2. If A and B are bounded context-free languages then is also a bounded context-free language. 3. If A is a bounded context-free language and x and y are fixed strings then the following set is also a bounded context-free language. ay Corollary 4.13. Every bounded context-free language is L-printable. Proof. Every finite set is L-printable. The L-printable sets are closed under the three properties in Theorem 4.12. Corollary 4.14. Every sparse context-free language is L-printable. This completely characterizes the L-printable context-free languages. However the sparse context-free languages do not characterize the L-printable languages. Proposition 4.15. There exists an L-printable set S such that S is not context-free Proof. The language is L-printable, but is not context-free. 5. L-Isomorphisms. It is easy to show that two P-printable sets, or P-rankable sets, of similar densities are P-isomorphic. Since the usual proof relies on binary search, it does not immediately extend to L-rankable sets. However, we are able to exploit the sparseness of L-printable sets to show the following. Theorem 5.1. If A and B are L-printable and have similar densities, then A and B are L-isomorphic (i.e., A - =log B). Proof. For each x, define y x to be the image of x in the lexicographic isomorphism from A to B. Since A and B are L-printable, they are both sparse. Let p(n) be a strictly increasing polynomial that bounds the densities of both sets. If x is "close" to y x , in the sense that there are at most p(jxj) strings between them in the lexicographic ordering. (Recall Definition 2.7.) In fact, for all x, jy x Let r A (x) be the rank of x in A. If A, then the rank of x in A is Furthermore, is the unique element of B for which this holds. Note that both r A (x) and r B (y x ) can be written in space O(log jxj). Thus, to compute y x , we need to compute do so by maintaining a variable d, that is initialized to r A (x). Counter c is initialized to 0. The following loop is iterated until a counter, c, reaches p(jxj 1. L-print (in lexicographic order) the elements of B of length c; for each string that is lexicographically smaller than decrement d; 2. increment c. Note that, if d is written on the work tape, each bit of x \Gamma d can be computed in logspace as needed, and the output of the L-printing function can be compared to in a bit-by-bit manner. L-PRIN If x 2 A, since the L-printing function outputs strings in lexicographic order, computing y x is easy: compute r A (x), then "L-print" B internally, actually outputting the r A (x) th string. Without loss of generality, we can assume that the simulated L-printer for B prints B in lexicographic order. Thus, as soon as the r A 1st element of B is printed internally, the simulation switches to output mode. The following is an overview of the logspace algorithm computing the desired isomorphism. 1. Compute A(x). 2. Compute r A (x), and write it on a work tape. 3. If x 2 A, find the r A (x) th element of B, and output it. 4. If find the unique string y and output y x . Using this theorem, we can now characterize the L-printable sets in terms of isomorphisms to tally sets, and in terms of sets of low Kolmogorov space complexity. Theorem 5.2. The following are equivalent: 1. S is L-printable. 2. S is L-isomorphic to some tally set in L. 3. There exists a constant k such that S ' KS[k log n; k log n] and S 2 L. Although it is not known whether or not every sparse L-rankable set is L-isomor- phic to a tally set (see Theorem 6.1), we can prove the following lemma, that will be of use in the proof of Theorem 5.2. Lemma 5.3. Let A be sparse and L-rankable. Then there exists a tally set T 2 L such that A and T have similar density. Proof. Let A -n denote the strings of length at most n in A. Let p(n) be an everywhere positive monotonic increasing polynomial such that jA -n j - p(n) for all n, and such that greater than the number of strings of length n in A. Let r(x) be the ranking function of A. We define the following tally set: To show that T 2 L, notice that of the tally strings 1 i , the largest n such that p(n \Gamma that n can be written in binary in space O(log m).) Then compute d This difference is bounded by p(n), and thus can be written in logspace. Finally, compute d compare to d 1 . Accept iff d 1 - d 2 . Finally, we show that T and A have similar density. Let f : A ! T be the lexicographic isomorphism between T and A. Note that f maps strings of length n to strings of length at most p(n), so f is polynomially bounded. Note that p is always positive, which implies that f is length-increasing. must also be polynomially bounded. Thus, T and A have similar density. The following proof of Theorem 5.2 is very similar to the proof of the analogous theorem in [AR88]. Proof. [1 be L-printable. Then it is sparse and L-rankable. Let T be the tally set guaranteed by Lemma 5.3. By Proposition 4.1, T is L-printable. Thus, T and S are L-printable, and T and S have similar density. So by Theorem 5.1, L. FORTNOW, J. GOLDSMITH, M. LEVY, AND S. MAHANEY [2 be L-isomorphic to a tally set T , and let f be the L-isomorphism from S to T . Let x 2 S be a string of length n. Let logspace computable, there exists a constant c such that r - n c , i.e., jrj - c log n. In order to recover x from r, we only have to compute f Computing 0 r given r requires log n space for one counter. Further, there exists a constant l such that computing requires at most lc log n space, since r - n c . So, the total space needed to compute x given r is less than or equal to log n+ lc log n - k log n for some k. Hence, log n]. If T 2 L, then S 2 L, since S - =log T . [3 log n] for some k, and S 2 L. On input 0 n , we simulate M u for each string of length k log n. For a given string x, log n, we first simulate M u (x) and check whether it completes in space k log n. If it does, we recompute M u (x), this time checking whether the output is in S. If it is, we recompute M u (x), and print out the result. The entire computation only needs O(log n) space, so S is L-printable. It was shown in [AR88] that a set has small generalized Kolmogorov complexity if and only it is P-isomorphic to a tally set. (Note: this was improvement of the result in [BB86], which showed that a set has small generalized Kolmogorov complexity if and only if it is "semi-isomorphic" to a tally set.) Using a similar argument and Theorem 5.2 we can show an analogous result for sets with small generalized Kolmogorov space complexity. First, we prove the following result. Proposition 5.4. For all M v and k, KS v [k log n; k log n] is L-printable. Proof. To L-print for length n, simulate M v on each string of length less than or equal to k log n, and output the results. Corollary 5.5. There exists a k such that A ' KS[k log n; k log n] if and only if A is L-isomorphic to a tally set. Proof. Suppose A is L-isomorphic to a tally set. Then, by the argument given in the proof of [2 ) 3] in Theorem 5.2, A ' KS[k log n; k log n]. suppose A ' KS[k log n; k log n]. By Proposition 5.4 and Theorem 5.2, KS[k log n; k log n] is L-isomorphic to a tally set in L via some L-isomorphism f . It is clear that A is L-isomorphic to f(A). Since f(A) is a subset of a tally set, f(A) must also be a tally set. 6. Printability, Rankability and Decision. In this section we examine the relationship among L-printable sets, L-rankable sets and L-decidable sets. We show that any collapse of these classes, even for sparse sets, is equivalent to some unlikely complexity class collapse. Theorem 6.1. The following are equivalent: 1. Every sparse L-rankable set is L-printable. 2. There are no tally sets in Proof. [2 , 3] This equivalence follows from techniques similar to those of Suppose A is a sparse L-rankable set. Note that A 2 L. Let ith bit of the jth string in A is 1g; where L-PRIN Note that hi; ji can be computed in space linear in jij + jjj. Since A is sparse, i and are bounded by a polynomial in the length of the jth string. Hence, hi; ji can be computed using logarithmic space with respect to the length of the jth string. Given hi; ji, we can determine i and j in polynomial time, and we can find the jth string of A by using binary search and the ranking function of A. Hence, T 2 P. So, by assumption, T 2 L. Next we give a method for printing A in logspace. Given a length n, we compute (and store) the ranks of 0 n and 1 n in A. Let r start and r end be the ranks of 0 n and the string with rank r start has length less than n. First, we check to see if 0 n 2 A, and if so, print it. Then, for each j, r start we output the jth string by computing and printing T (1 hi;ji ) for each bit i. This procedure prints the strings of A of length n. Note that since A is sparse, we can store r start and r end in O(log n) space. Since store and increment the current value of i in log n space. [1 be a tally set. Since the monotone circuit value problem is P-complete (see [GHR95]), there exists a logspace-computable function f and a nondecreasing polynomial p such that f(n) produces a circuit Cn with the following properties. 1. Cn is monotone (i.e., Cn uses only AND and OR gates). 2. Cn has p(n) gates. 3. The only inputs to Cn are 0 and 1. 4. Cn outputs 1 iff 1 n is in T . We can assume that the reduction orders the gates of Cn so that the value of gate depends only on the constants 0 and 1 and the values of gates g j for ([GHR95]). Let xn be the string of length p(n) such that the ith bit of xn is the value of gate g i . Ng. Then A contains exactly one string of length p(n) for all n, and no strings of any other lengths. 6.1.1. The set A is L-rankable. Proof. To prove this claim, let w be any string. In logspace, we can find the greatest n such that p(n) - jwj. If p(n) 6= jwj then w 62 A, and the rank of w is n. Suppose xn is the only string of length p(n) in A, the rank of w is Consider the ith bit of w as a potential value for gate g i in Cn . Let j be the smallest value such that w j is not the value of g j . In order to find the value of a gate i , we first use f(n) (our original reduction) to determine the inputs to g i . By the time we consider the i th bit of w, we know that w is a correct encoding of all of the gates g k such that k ! i, so we can use those bits of w as the values for the gates. Thus, we can determine the value of g i and compare it to the ith bit of w. If they differ, we are done. If they are the same, we continue with the next gate. We can count up to p(n) in logspace, so this whole process needs only O(log p(n)) space to compute. Once j is found, there are three cases to consider. 1. If j doesn't exist then 2. If the jth bit of w is 0 then w ! xn . 3. If the jth bit of w is 1 then w ? xn . These follow since the ith bit of xn matches the ith bit of w for all Thus A is L-rankable and, by assumption, L-printable. 12 L. FORTNOW, J. GOLDSMITH, M. LEVY, AND S. MAHANEY So, to determine if 1 n is in T , L-print A for length p(n) to get xn . The bit of xn that encodes the output gate of Cn is 1 iff 1 every step of this algorithm is computable in logspace, T 2 L. This completes the proof of Theorem 6.1. Corollary 6.2. There exist two non-L-isomorphic L-rankable sets of the same density, unless there are no tally sets in Proof. Consider the sets T and A from the second part of the proof of Theorem 6.1. The set has the same density as A. By Proposition 4.1, B is L-printable. If A and B were L-isomorphic then by Proposition 3.6, A would also be L-printable and T would be in L. One may wonder whether every sparse set in L is L-printable or at least L- rankable. We show that either case would lead to the unlikely collapse of FewP and L. Recall that FewP consists of the languages in NP accepted by nondeterministic polynomial-time Turing machines with at most a polynomial number of accepting paths. Fix a nondeterministic Turing machine M and an input x. Let p specify an accepting path of M(x) represented as a list of configurations of each computation step along that path. Note that in logarithmic space we can verify whether p is such an accepting computation since if one configuration follows another only a constant number of bits of the configuration change. We can assume without loss of generality that all paths have the same length and that no accepting path consists of all zero or all ones. Define the set PM by is an accepting path of M on xg: From the above discussion we have the following proposition that we will use in the proofs of Theorems 6.6 and 6.7. Proposition 6.3. For any nondeterministic machine M , PM is in L. Allender and Rubinstein [AR88] showed the following about P-printable sets. Theorem 6.4 ([AR88]). Every sparse set in P is P-printable if and only if there are no sparse sets in FewP \Gamma P. Allender [All86] also relates this question to inverting functions. Definition 6.5. A function f is strongly L-invertible on a set S if there exists a logspace computable function g such that for every x 2 S, g(x) prints out all of the strings y such that We extend the techniques of Allender [All86] and Allender and Rubinstein [AR88] to show the following. Theorem 6.6. The following are equivalent. 1. There are no sparse sets in FewP \Gamma L. 2. Every sparse set in L is L-printable. 3. Every sparse set in L is L-rankable. 4. Every L-computable, polynomial-to-one, length-preserving function is strongly L-invertible on f1g . 5. Proof. [1 A be a sparse set in L. Then A is in P. By (1) we have that there are no sparse sets in FewP \Gamma P. By Theorem 6.4, A is P-printable. Consider the following set B. ith bit of the jth element of A of length n is bg L-PRIN Since A is P-printable then B is in P. By (1) (as B is sparse and in P ' FewP), we have B is in L. Then A is L-printable by reading the bits off from B. [2 ) 3] Follows immediately from Proposition 3.4. [3 A be a sparse set in FewP accepted by a nondeterministic machine M with computation paths of length q(n) for inputs of length n. Consider the set PM defined as above. Note that PM is sparse since for any length accepts a polynomial number of strings with at most a polynomial number of accepting paths each. Also by Proposition 6.3 we have PM in L. By (3) we have that PM is L-rankable. We can then determine in logarithmic space whether M(x) accepts (and thus x is in A) by checking whether r PM (x#0 [2 f be a L-computable, polynomial-to-one, length-preserving function. Consider g. Since S is in L, S is L-printable. A be a sparse set in L. Define x is in A and x otherwise. If g is a strong L-inverse of f on 1 then g(1 n ) will print out the strings of length n of A and 1 n . We can then print out the strings of length n in logspace by printing the strings output by g(1 n ), except we print 1 n only if 1 n is in A. [1 , 5] In [RRW94], Rao et al. show that there are no sparse sets in FewP \Gamma P if and only if straightforward modification of their proofs is sufficient to show that there are no sparse sets in FewP \Gamma L if and only if Unlike L-printability, L-rankability does not imply sparseness. One may ask whether every set computable in logarithmic space may be rankable. We show this equivalent to the extremely unlikely collapse of PP and L. Theorem 6.7. The following are equivalent. 1. Every #P function is computable in logarithmic space. 2. 3. Every set in L is L-rankable. Our proof uses ideas from Blum (see [GS91]), who shows that every set in P is P-rankable if and only if every #P function is computable in polynomial time. Note that Hemachandra and Rudich [HR90] proved results similar to Blum's. Proof. [1 A is in PP then there is a #P function f such that x is in A iff the high-order bit of f(x) is one. [2 implies that . Thus we have and we can compute every bit of a #P function in logarithmic space. [1 A be in L. Consider the nondeterministic polynomial-time machine M that on input x guesses a y - lex x and accepts if y is in A. The number of accepting paths of M(x) is a #P function equal to r A (x). [3 f be a #P function. Let M be a nondeterministic polynomial-time machine such that f(x) is the number of accepting computations of M(x). Let q(n) be the polynomial-sized bound on the length of the computation paths of M . Consider PM as defined above. By Proposition 6.3 we have that PM is in L so by (3) PM is L-rankable. We then can compute f(x) in logarithmic space by noticing 14 L. FORTNOW, J. GOLDSMITH, M. LEVY, AND S. MAHANEY 7. Conclusions. The class of L-printable sets has many properties analogous to its polynomial-time counterpart. For example, even without the ability to do binary searching, one can show two L-printable sets of the same density are isomorphic. However, some properties do not appear to carry over: it is very unlikely that every sparse L-rankable set is L-printable. Despite the strict computational limits on L-printable, this class still has some bite: every tally set in L, every sparse regular and context-free language and every L-computable set of low space-bounded Kolmogorov complexity strings is L-printable. Acknowledgments . The authors want to thank David Mix Barrington for a counter-example to a conjecture about sparse regular sets, Alan Selman for suggesting the tally set characterization of L-printable sets and Corollary 5.5, Chris Lusena for proofreading, and Amy Levy, John Rogers and Duke Whang for helpful discussions. The simple proof sketch of Theorem 4.3 was provided by an anonymous referee. The last equivalence of Theorem 6.6 was suggested by another anonymous referee. The authors would like to thank both referees for many helpful suggestions and comments. --R The complexity of sparse sets in P Sets with small generalized Kolmogorov complexity Structural Complexity I Ambiguity in graphs and expressions Tally languages and complexity classes The Mathematical Theory of Context-Free Languages Limits to Parallel Computation: P-Completeness Theory Compression and ranking Generalized Kolmogorov complexity and the structure of feasible com- putations On sparse oracles separating feasible complexity classes Computation times of NP sets of different densities ambiguity and other decision problems for acceptors and transducers Alternierung und Logarithmischer Platz Introduction to Languages and the Theory of Computation Computational Complexity Upward separation for FewP and related classes A note on sets with small generalized Kolmogorov complexity A complexity theoretic approach to randomness Characterizing regular languages with polynomial densities --TR --CTR Allender, NL-printable sets and nondeterministic Kolmogorov complexity, Theoretical Computer Science, v.355 n.2, p.127-138, 11 April 2006

context-free languages;sparse sets;kolmogorov complexity;ranking;l-isomorphisms;computational complexity;regular languages;logspace

298506

The Inverse Satisfiability Problem.

We study the complexity of telling whether a set of bit-vectors represents the set of all satisfying truth assignments of a Boolean expression of a certain type. We show that the problem is coNP-complete when the expression is required to be in conjunctive normal form with three literals per clause (3CNF). We also prove a dichotomy theorem analogous to the classical one by Schaefer, stating that, unless P=NP, the problem can be solved in polynomial time if and only if the clauses allowed are all Horn, or all anti-Horn, or all 2CNF, or all equivalent to equations modulo two.

Introduction . Logic deals with logical formulae, and more particularly with the syntax and the semantics of such formulae, as well as with the interplay between these two aspects [CK90]. In the domain of Boolean logic, for example, a Boolean formula OE may come in a variety of syntactic classes-conjunctive normal form (CNF), its subclasses 3CNF, 2CNF, Horn, etc.- and its semantics is captured by its models or satisfying truth assignments, that is, the set -(OE) of all truth assignments that satisfy the formula (see Figure 1 for an example). Going back and forth between these two representations of a formula is therefore of interest. One direction has been studied extensively from the standpoint of computational complexity: going from OE to -(OE). In particular, telling whether is the famous satisfiability problem (SAT), which is known to be NP-complete in its generality and its special case 3SAT, among others, and polynomial-time solvable in its special cases Horn, 2SAT, and exclusive-or [Co71, Sc78, Pa94]. All in all, this direction is a much-studied computational problem. In this paper we study, and in a certain sense completely settle, the complexity of the inverse problem, that is, going from -(OE) back to OE. That is, for all the syntactic classes mentioned above, we identify the complexity of telling, given a set M of models, whether there is a formula OE in the class (3SAT, Horn, etc.) such that We call this problem inverse satisfiability. Besides its fundamental nature, there are many more factors that make inverse satisfiability a most interesting problem. A major motivation comes from AI (in fact, what we call here the inverse satisfiability problem is implicit in much of the recent AI literature [Ca93, DP92, KKS95, KKS93, KPS93]). A set of models such as those in Figure 1(b) can be seen as a state of knowledge. That is, it may mean that at present, for all we know, the state of our three-variable world can be in any one of the three states indicated. In this context, formula OE is some kind of knowledge representation. In AI there are many sophisticated competing methods for knowledge representation Received by the editors April 24, 1995; accepted for publication (in revised form) November 20, 1996; published electronically June 15, 1998. This work was partially supported by the Esprit Project ALCOM II and the Greek Ministry of Research (\PiENE\Delta program 91E\Delta648). http://www.siam.org/journals/sicomp/28-1/28511.html y Department of Mathematics, University of Patras, Patras, Greece ([email protected]). z Department of Computer Science, Athens University of Economics and Business, Athens, Greece ([email protected]). (a) (b) Fig. 1. A Boolean formula in 3CNF (a), and the corresponding set of models (b). (Boolean logic is perhaps the most primitive; see [GN87, Le86, Mc80, Mo84, Re80, SK90]), and it is important to understand the expressibility of each. This is a form of the inverse satisfiability problem. The inverse satisfiability problem was also proposed in [DP92] as a form of discovering structure in data. For example, establishing that a complex binary relation is the set of models of a simple formula may indeed uncover the true structure and nature of the heretofore meaningless table. [DP92] only address this problem in certain fairly straightforward cases. The problem of learning a formula [AFP92] can be seen as a generalization of the inverse satisfiability problem. A recent trend in AI is to approximate complex formulae by simple ones, such as Horn formulae [SK91, KPS93, GPS94]. Quantifying the quality and computational feasibility of such approximations also involves understanding the inverse satisfiability problem. The basic computational problem we study is this: given a set of models M , is there a CNF formula OE with at most three literals per clause, such that We call this problem INVERSE 3SAT. Our first result is that INVERSE 3SAT is coNP-complete (Theorem 1). Note. INVERSE 3SAT, as well as all other problems we consider in this paper, can be solved in polynomial time if the given m \Theta n table M has if there are exponentially many models in M . The interesting cases of the problem are therefore when There are three well-known tractable cases of SAT: 2SAT (all clauses have two literals), HORNSAT (all clauses are Horn, with at most one positive literal each, and its symmetric case of anti-Horn formulae, in which all clauses have at most one negative literal), and XORSAT (the clauses are equations modulo two). Schaefer's elegant dichotomy theorem [Sc78] states that, unless P=NP, in a certain sense these are precisely the only tractable cases of SAT. Interestingly, the inverse problem for these three cases happens to also be tractable! That is, we can tell in polynomial time if a set of models is the set of models of a Horn (or anti-Horn) formula, of a 2CNF formula, or of an exclusive-or formula (interestingly, the latter two results were in fact pointed out by Schaefer himself [Sc78], while the first, left open in [Sc78], is from [DP92, KPS93]). The question comes to mind: are there other tractable cases of the inverse problem? Our Theorem 2 answers this in the negative; rather surprisingly, a strong dichotomy theorem similar to Schaefer's holds for the inverse satisfiability problem as well, in that the problem is coNP-complete for all syntactic classes of CNF formulae except for the cases of Horn (and anti-Horn), 2CNF, and exclusive-or. The proof of our dichotomy theorem draws from both that of Theorem 1 and Schaefer's proof, and in fact strengthens Schaefer's main expressibility result (Theorem 3.0 in [Sc78]). 154 DIMITRIS KAVVADIAS AND MARTHA SIDERI 2. Definitions. Most of the nonstandard terminology used in this paper comes from [Sc78]. be a set of Boolean variables. A literal is a variable or its negation. A model is a vector in f0; 1g n , intuitively a truth assignment to the Boolean variables. We denote by - and - the logical or and and, respectively. We also extend this notation to bitwise operations between models. If t is a model, we denote by t i the constant (i.e., 0 or 1) in the ith position of t. A k-place logical relation is a subset of f0; 1g k (k integer). We use the notation [OE], where OE is a Boolean formula, to denote the relation defined by OE when the variables are taken in lexicographic order. Let R be a logical relation. Call R Horn if it is logically equivalent to a conjunction of clauses, each with at most one positive literal. We call it anti-Horn if it is equivalent to a conjunction of clauses with at most one negative literal. We call it 2CNF if it is equivalent to a 2CNF expression. Finally, we call it affine if it is the solution of a system of equations in the two-element field. be a set of Boolean relations. An S-clause (of arity is an expression of the form R(a is a k-ary relation in S and the a i 's are either Boolean literals or constants (0 or 1). Given a truth assignment, we consider an S-clause to be true if the combination of the constants, if any, and the values assigned to the variables form a tuple in R. Define an S-formula to be any conjuction of S-clauses defined by the relations in S. The generalized satisfiability problem is the problem of deciding whether a given S-formula is satisfiable. Schaefer's dichotomy theorem [Sc78] states that the satisfiability of an S-formula can be decided in polynomial time in each of the following cases: (a) all relations in S are Horn, (b) all relations in S are anti-Horn, (c) all relations in S are 2CNF, (d) all relations in S are affine. In all other cases the problem is NP-complete. That is, Schaefer's result totally characterizes the complexity of the CNF satisfiability problem where in addition, the clauses are allowed to be arbitrary relations of bounded arity. It is interesting to note that several restricted forms of SAT such as ONE-IN-THREE 3SAT, NOT-ALL-EQUAL 3SAT etc., all follow as special cases of generalized satisfiability (see [GJ79, Pa94]). To make this point more clear, notice that the problem ONE-IN-THREE 3SAT can be considered as a set of four 3-ary relations g. The first relation is and corresponds to the S-clause R 1 relation is ff0; 0; 0g; f1; 1; 0g; f1; 0; 1gg and corresponds to the S-clauses with one negated literal, e.g., R 2 For any Boolean formula OE we denote by -(OE) its set of models. We say that a set of models M is a 3CNF set (kCNF in general) if there is a formula OE in 3CNF (respectively, kCNF) such that -(OE). Notice that for any model set M we can construct a kCNF formula that has M as its model set, but in general, this may require extra existentially quantified variables. Based on the above we define the INVERSE SAT problem for a set of relations S as follows. Given a set M ' f0; 1g n , is there a conjunction of S-clauses over n variables that has M as its set of models? Our main result states that if the relations fall in each of the four cases above, the INVERSE SAT problem is also polynomial. Otherwise it is coNP-complete. Notice that we have excluded S from being part of the instance since we want to emphasize that INVERSE SAT is actually a collection of infinitely many subproblems. This means that all relations of S are of constant arity. Otherwise, relations of non- constant arity could have exponentially many tuples and the problem becomes trivially intractable. In the next section we prove that the INVERSE SAT problem is coNP-complete for 3CNF formulas. This proof includes the main construction that will be used in the proof of the main theorem in the last section. This last proof makes use of an expressibility result which is interesting on its own and partially relies on Schaefer's main theorem but with several interesting extensions. 3. coNP-completeness of inverse 3SAT. We begin this section with a technical definition that will be used throughout the paper. Definition. Let n be a positive integer and let M ' f0; 1g n be a set of Boolean vectors. For k ? 1, we say that a Boolean vector m 2 f0; 1g n is k-compatible with M if for any sequence of k positions a vector in M that agrees with m in these k positions. The above definition implies that a vector m 2 f0; 1g n is not k-compatible with a set of Boolean vectors M if there exists a sequence of k positions in m that does not agree with any vector of M . The following is a useful characterization of kCNF sets. Lemma 1. Let M ' f0; 1g n be a set of models. Then the following are equivalent. (a) M is a kCNF set. Proof. Let OE M be the conjuction of all possible kCNF clauses defined on n variables and satisfied by all models in M . Notice that OE M is the most restricted kCNF formula (in terms of its model set) which is satisfied by all models in M . Hence if (a) holds, does not satisfy at least one clause of OE M and concequently disagrees with all models in M in the same k positions corresponding to the variables in the clause, that is, m is not k-compatible with M . Conversely, assume that any model not in M is not k-compatible with M . Then means m does not satisfy OE differs from all members of M in some k positions, so the k-clause indicating the complement of m in those k positions is in does not satisfy OE M . So is a kCNF set. The INVERSE 3SAT problem is this: given a set of models M , is it a 3CNF set? We now state our first complexity result. Theorem 1. INVERSE 3SAT is coNP-complete. Proof. Lemma 1 establishes that the problem is in coNP: given a set M of models, in order to prove that it is not a 3CNF set, it suffices to produce a model that is 3-compatible with M (obviously, 3-compatibility can be checked in polynomial time). Alternatively, given M , we immediately have a candidate 3CNF formula OE the conjunction of all 3CNF clauses that are satisfied by all models in M . Thus M is not a 3CNF set iff there is a model not in M that satisfies OE M . To prove coNP-completeness, we shall reduce the following well-known coNP-complete problem to INVERSE 3SAT: given a 3CNF formula, is it unsatisfiable? Given a 3CNF formula / with n - 4 variables and c clauses, we shall construct a set of models M such that M is 3CNF iff / is unsatisfiable. The set M will contain models, one for each set W of three variables, and each truth assignment T to these three variables that does not contradict a clause of / (since we may assume that / consists of clauses that have exactly three literals each). Let W be a set of three variables chosen among the variables fx of formula /, and let T : W 7! f0; 1g be a truth assignment to the variables of W , such that / does not contain a clause not satisfied by T . Consider some total order among the pairs (W; T ), say the lexicographic one. The set M will contain a model mW;T for each W and T and no other model. Every boolean vector mW;T is a concatenation of the encodings - T occuring in the formula /. The encoding - T of a variable x is a Boolean vector of length and is defined as follows: positions z - where (W; T ) is the ith pair in the total order mentioned above. Notice that if x 2 W , the value of x in T is determined by the first two positions of - T the code 01 stands for the value 1, and the code 10 stands for x being 0. In these two cases we call the string - T W (x) a value pattern. When x 62 W , the code 00 in the first two positions denotes the absence of x from W , while the rest of the string uniquely determines the pair (W; T ). In this case we call the string - T W (x) a padding pattern. Notice that by our construction in a vector are exactly occurences of the unique padding pattern for (W; T ), while the remaining three are value patterns. Hence, the length of each Boolean vector mW;T is n(k there is no exponential blow-up in the construction of the set M . The proof of Theorem 1 now rests on the next claim. Claim. There is a model not in M that is 3-compatible with M if and only if / is satisfiable. Proof of the claim. For the moment, consider a Boolean vector where the length of each substring m i 2. It is obvious that if the model m is 3-compatible with M , then it is 3-compatible in the positions restricted to one substring That is, if we take three arbitrary positions of m i , there is a vector mW;T in M that agrees with m i in these three positions. The 3- compatibility of m i with M also implies something stronger: that there is a vector which contains a substring - T To see this, first assume that m i does not have the value 1 in any position j for 3 2. Then 3-compatibility forces m i to have the values 0 and 1 or 1 and 0 in the first and second positions; i.e., m i is a value pattern. Now, if m i has the values 0 and 1 in positions 2, then the values in any triple of positions that includes positions can only agree with the values in the same positions of a specific model of M , namely, the one having the padding pattern with 0 in position and 1 in position j. Therefore, m i is identical to this padding pattern. In this case, however, an analogous observation shows that the whole 3-compatible model m is identical to the model of M that has this pattern. So if m is 3-compatible with M , either it is in M or it consists of value patterns only. Assume now that there exists a model 2 M that is 3-compatible with M . As already proved, this model consists only of value patterns m i . Model m encodes a satisfying truth assignment to the variables of /. For suppose it did conflict with a clause c of / over variables fx i g. Consider the three value patterns in the positions of the variables of c. Since m is 3-compatible with M and each value pattern contains only one 1, we can conclude that there exists a model which encodes a truth assignment T to the set of variables such that - T But since by construction T does not contradict a clause of /, we couldn't have conflicted with a clause of /. Therefore, the Boolean vector is an encoding of a satisfying assignment to the variables for formula /: string m i is an encoding of the truth value assigned to the variable x i for each formula / is satisfiable since every clause of / is satisfied by the truth assignment described by vector m. Conversely, assume that / is satisfiable; i.e., there exists a satisfying truth assignment s for the variables fx g. Construct the model concatenation of value patterns, where every string m i is defined as follows: positions z - Obviously, model m is not included in the set M , since every model in M contains a padding pattern. Suppose that m is not 3-compatible with M . In this case m contains three positions that do not agree with any model in M . Since m is a concatenation of value patterns, it must contain three substrings assignment T for the set of variables such that the pair (W; T ) is not encoded in any model of M . All sets of variables are, however, examined during the construction of M , and the only truth assignments that are not encoded are those conflicting with a clause of /. Since T does not conflict with any clause-because it is a restriction of s to three variables-we conclude that the pair (W; T ) is encoded in some model of M . Hence, m is 3-compatible with M . So, if / is satisfiable, there exists a model 3-compatible with M , specifically the model encoding a satisfying truth assignment. 4. The dichotomy theorem. Our main result is the following generalization of Theorem 1. Theorem 2. The INVERSE SAT problem for S is in PTIME in each of the following cases. (a) All relations in S are Horn. (b) All relations in S are anti-Horn. (c) All relations in S are 2CNF. (d) All relations in S are affine. In all other cases, the INVERSE SAT problem for S is coNP-complete. [Sc78] proves a surprisingly similar dichotomy theorem for SAT: SAT is in PTIME for all of these four classes, and NP-complete otherwise. Our proof is based on an interesting extension of Schaefer's main result, explained below. Definition. Let S be a set of Boolean relations and let R be another Boolean relation, of arity r. We say that S faithfully represents R if there are binary Boolean functions s such that there is a conjunction of S-clauses over the variables which is logically equivalent to the formula s for some That is, S-clauses can express R with the help of uniquely defined auxiliary variables. This is a substantial restriction of Schaefer's notion of ``represents,'' which allows arbitrary existentially quantified conjunctions of S-clauses (our definition only allows quantifiers which are logically equivalent to 9!x). Hence our main technical result below extends the main result of [Sc78, Theorem 3.0]. Independently, Creignou and Hermann [CH96] have defined the concept "quasi-equivalent," which is the same as the concept of "faithful representation" defined in this paper. Theorem 3. If S does not satisfy any of the four conditions of Theorem 2; then faithfully represents all Boolean relations. Proof. Assuming that none of the four conditions are satisfied by S, the proof proceeds by finding more and more elaborate Boolean relations that are faithfully represented by S. Notice that, since the notion of faithful representation was defined as equivalence of two S-formulas, we shall restrict the proof to the construction of appropriate S-clauses-faithful representation of the corresponding relations will then follow immediately. In this process the allowed operations must preserve the uniqueness of the values of the auxiliary variables and produce a formula which is also in conjuctive form. Therefore, if C and C 0 are S-formulas, the allowed operations are: (a) C - C 0 , i.e., conjuction of two S-formulas, (b) C[a=x], i.e., substitution of a variable symbol by another symbol, (c) C[0=x] and C[1=x], i.e., substitution of a variable by a constant (this is actually a selection of the tuples that agree in the specified constant), and (d) 9!xC(x), i.e., existential quantification, where the bound variables are uniquely defined. Some of the steps are provided by Schaefer's proof, and some are new. Step 1. Expressing [x j y]. This was shown in [Sc78, Lemma 3.2 and Corollary 3.2.1]. The following exposition is somewhat simpler and is based on the fact that a set M ' f0; 1g n is the model set of a Horn formula iff it is closed under bitwise -; see the Appendix and [KPS93]. Let R be any non-Horn relation of S (say of arity k). The closure property mentioned above implies that there exist models t and t 0 in R such that t - t Based on R we may define the clause R set a all positions i where both t i and t i 0 are 0 (resp., 1). Set a x to all positions where y to all positions where t 1. It is easy to see that both x and y actually appear in R 0 . (If not, then one of t and t 0 coincides with their conjunction.) Now 01 and 10 are models of R 0 , but 00 is not. Hence R 0 is either in addition, S contains a relation which is not anti-Horn, then a symmetric argument rules out tuple 11, resulting in a clause R 00 which is either Notice that since this is the case we shall henceforth feel free to use negative literals in our expressions. Step 2. Expressing [x -y]. Schaefer shows in Lemma 3.3 that there is an S-clause involving variables x; y; z whose set of models contains 000; 101; 011, but not 110. The proof is as follows: it is known (see the Appendix) that an S-clause is affine if and only if for any three models t is also a model. Consider, therefore, an S-clause that is not affine and assume that [x j y] can be represented. By the observation in Step 1 we may negate the variables of the clause in the positions where t 0 is 1. Now the new S-clause, call it S 0 , is satisfied by the all-zero truth assignment and moreover by the assignments t 1 not by 0 \Phi t 1 . Construct a new clause R(a 1 ; a is the arity of positions i where both t 1 are 0, a both are 1, a 0 is 1, and finally a 0 is 0. The S-clause R defined on x; y; z, has models 000, 011, 101 (corresponding to the all-zero assignment, t 1 0 of S 0 , respectively), but not 110 (which corresponds to t 1 We will show that R faithfully represents one of the four versions of or: Observe that at least two of x; actually occur in R. If exactly two variables are present in R, then R represents a version of or as follows: if x and y are present, then R(x; are present, then are present, then R(y; z). If all three variables are present, depending on which of the remaining four possible models are also in the model set of the S-clause, we have sixteen possible relations. Of these, the strongest, with models identical to the set can be used to define X(x; (which is true when exactly one of x; y; z is true) as follows: X(x; and in this case the current step is unnecessary. In each of the other fifteen cases, we show by exhaustive analysis that there is an R-clause, with one constant, which represents a version of or. If then R(0; then R(x; 0; Since we can also faithfully express [w j x], by Step 1, we have all four versions of or. Step 3. Expressing X(x; z). X is a formula which is satisfied if exactly one of the three variables has the value 1. It is known (see the Appendix) that an S- clause is 2CNF iff for any set of three satisfying assignments t 0 , t 1 , t 2 , the assignment satisfying assignment. We use this characterization to prove that if a relation set S contains a relation which is not 2CNF and also contains relations which are not Horn, anti-Horn, and affine, then X(x; y; z) can be faithfully represented. Consider an S-clause which is not 2CNF. We may therefore find three satisfying assignments such that the expression (t a satisfying assignment. As in the previous step we may negate the variables in the positions where t 0 has the value 1, resulting in a new clause S 0 , which is satisfied by the all-zero assignment, by t 1 not by t 1 which is equal to t 1 0 . Set 0 to all positions where both t 1 are 0, x to all positions where both t 1 are 1, y where t 1 0 is 1, and finally z where 0 is 0. Observe that all three variables actually occur in the constructed clause R: if x is not present then t 1 0 is identical to the all-zero assignment, a contradiction; if either y or z is not present then t 1 0 is identical to t 1 again a contradiction. The clause constructed includes models 000, 110, and 101, but not 100. Now the S-clause R(x; exactly the models 100, 010, and 001; i.e., it is X(x; Step 4. Expressing [x j (y-z)]. Notice that the expression X(x; s; y)-X(x; t; z)- X(s; t; u) is equivalent to Thus we prove that we can faithfully represent a relation in which a variable is logically equivalent to the or of two other variables. Notice that the auxiliary variables s; t; u are uniquely defined by the values of y and z. Step 5. Using repeatedly [x j (y - z)] and [x j y] we can faithfully represent any clause, and by taking conjunctions of arbitrary clauses we can faithfully represent any Boolean relation, completing the proof of Theorem 3. Proof of Theorem 2. Let S be a set of relations satisfying one of conditions (a)- (d), and let r be the maximum arity of any relation in S; we can solve the inverse satisfiability problem for S as follows. Given a set of models M , we first identify in time O(n r jM all S-clauses that are satisfied by all models in M ; call the conjunction of these S-clauses OE. Clearly, if there is a conjunction of S-clauses that has M as its set of models, then by the arguments used in Lemma 1, it is precisely OE. To tell whether the set of models of OE is indeed M , we show how to generate the set of models of OE with polynomial delay between consecutive outputs [JPY88]. Provided that such generation is possible, we can decide whether checking if the generated models belong in M . If a model not in M is generated, then we reply "no"; otherwise, if the set of models generated is exactly M , we reply "yes." Observe that the answer will be obtained after at most jM in overall polynomial time. Our generation algorithm is based on a more general observation that also explains the analogy of our dichotomy theorem to the one of Schaefer's. Call a syntactic form of a Boolean formula hereditary if the substitution of a variable by a constant results in a new formula of the same syntactic form. Observe that the four cases for which we claim that the inverse satisfiability problem is polynomial are indeed hereditary and coincide with the polynomial cases of satisfiability [Sc78]. Theorem 4. If the following two conditions hold for a class of Boolean formulas: (a) the syntactic form of the class is hereditary, and (b) the satisfiability problem for the class is in PTIME, then the models of any formula in the class can be generated with polynomial delay between consecutive outputs. Proof. Here is an informal description of the generation algorithm: at each step we substitute a variable by a constant, first by the value 1 and then by 0. Since (a) holds, the substitution results in a new formula of the same syntactic form. We then ask a polynomial-time oracle whether the produced formula is satisfiable. Since (b) holds, such an oracle exists. If the produced formula is satisfiable, we proceed recursively and substitute the next variable until all variables have been assigned a value, in which case we return the model. When at a certain step we are through with the value 1 for a variable (either by discovering a model or by rejecting the value because the produced formula is unsatisfiable), we try the value 0, and when finished, we backtrack to the previous step. It is easy to see that after at most 2n queries to the oracle (where n is the number of variables) we either generate a new model or we know that all models of the formula have been generated. Now, to show coNP-completeness of all other cases, let S be a set of Boolean relations not satisfying conditions (a)-(d). It is clear that the INVERSE SAT problem for S is in coNP: let r - 3 be the largest arity of any relation in S. Given a set of models M , we construct all S-clauses satisfied by all models in M-this takes time. M is the set of models of a conjunction of S-clauses if and only if all models not in M fail to satisfy at least one of these S-clauses. To show completeness, we shall reduce UNSATISFIABILITY, the problem of telling whether a 3CNF expression / is unsatisfiable, to the INVERSE SAT(S). We suppose that / is a 3CNF expression on n ? 3r variables. Set M contains a model for each 3r-tuple of variables and values for these variables that don't contradict any clause of /. Let k be the cardinality of M , a quantity bounded by a function of r and of the number of variables and clauses of /. Notice that since r is constant, the number of models is not exponential. Our construction is a generalization of that of Theorem 1: we consider some total order among the pairs (W; T ), where W is a set of 3r variables and T a truth assignment to those variables that does not contradict any clause of /. Every Boolean vector mW;T in M is a concatenation of two strings: W . String fi T W is a concatenation of the encodings - T occuring in the formula /: fi T The encoding of - T W (x) of a variable x is a Boolean vector of length and is defined in the proof of Theorem 1. Notice that in this construction the unique padding pattern for (W; T ) occurs in the string fi T W . Call the length of a string fi T W . The string ffl T W is constructed as follows: we consider all 3CNF clauses on N variables satisfied by the set of strings fi T W for all sets of 3r variables W and assignments T to those variables. Call OE the conjuction of all these clauses. We express OE faithfully by S-clauses. This will involve auxiliary variables . From the definition of faithful representation we see that xN+' j f ' however, that each of the auxiliary variables depends on at most three of the N variables appearing in the 3CNF clauses. This follows from the fact that we are representing 3CNF clauses, and consequently, we can express each 3CNF clause separately by S-clauses and then take the conjunction of the representations. Thus, the overall dependency of an auxiliary variable xN+' a Boolean function f ' be a string fi T W . The values in the s positions of the corresponding string ffl T are the values of the auxiliary variables: b (Note that these values are stated explicitly, i.e., not encoded as value patterns.) This is where the concept of faithful representation is necessary: for each string fi T W there is a unique string ffl T W . With ordinary representation the multiple ways to extend a string fi T W via the auxiliary variables would result in an exponential increase of our model set. Let M ' f0; 1g N+s be the constructed set of models. We claim that M is the set of models of a conjunction of S-clauses iff the original 3CNF expression / is unsatisfiable. If / is satisfiable, then M is not the set of models of any rCNF expression. Consider the model corresponding to the satisfying truth assignment. This model is a concatenation of two parts: the first has N positions and consists of the value patterns encoding the values of all variables in the satisfying truth assignment, exactly as in the proof of Theorem 1, and the second consists of the corresponding values of the s auxiliary variables. This model is r-compatible with M : any r-tuple restricted to the first N positions certainly matches a corresponding tuple in some model, by the construction of M . In fact, when the tuple is restricted to the first part, any 3r-tuple can be matched. This is precisely why an r-tuple that is not restricted to the first N positions is also r-compatible: by the dependency of each auxiliary value to at most 3 of the first N , a compatibility of an i-tuple (i - r) in the second part holds if a 3i-compatibility in the first part holds. Alternatively, instead of looking at a position in the second part, we can look at the three corresponding positions of the first part. Therefore, the whole model corresponding to the satisfying truth assignment is r-compatible with M . It follows by Lemma 1 that M is not rCNF, and as a result, M is not the set of models of any conjunction of S-clauses (recall that the maximum arity in S is r). Suppose then that / is unsatisfiable. Let M 0 be M restricted to the first N positions. Then M 0 is exactly the set of models of OE (the conjuction of all 3CNF clauses on N variables which don't disagree with M 0 ) by the reasoning in Theorem 1: no model is 3-compatible with M 0 except those in M 0 . Since M 0 is the set of models of OE, it follows that M is the set of models of the corresponding conjunction of S-expressions that faithfully represents OE. Appendix . This appendix contains the proof of the closure properties of Horn, anti-Horn 2CNF, and affine sets of models, which are used in the proof of Theorem 3. In what follows, M ' f0; 1g n denotes a set of models. Sets. M is Horn iff for any two models t; t the model (t - t 0 ) is also in M . The proof is based on the following proposition from [KPS93]. If t and t 0 are bit-vectors we use the notation t - t 0 to denote that t Proposition. The following are equivalent. (a) There is a Horn formula whose model set is M . (b) For each either there is no t or there is a unique (c) If t; t Proof. That (a) implies (c) is easy. To establish (b) from (c), take t 0 to be the - of all t 00 2 M such that t - t 00 . Finally, if we have property (b), we can construct the following set of Horn clauses: for each be the model guaranteed by (b); create a Horn clause (( 1. It is easy to see that the set of all these Horn clauses comprises the desired OE. Anti-Horn Sets. This case is symmetric to the above. Just replace 1 with 0 and - with -. Sets. M is 2CNF iff for any set of three models t the model Proof. This was shown in [Sc78, Lemma 3.1B]. We give a different proof, which is simpler and is based on Lemma 1 for 2. First notice that the model has the following property. The value of t in each position is equal to a value, which is the majority among the three values of the models position (e.g., if the values of models t position i are (1; 1; 0), respectively, the value of t in position i is 1). Call the outcome of the operation the majority model of t Only if: Suppose M is 2CNF. By Lemma 1 any 2-compatible model with M is in M . It is easy to see that the majority model of any three models is 2-compatible with these three models. If: Suppose that the majority model of any set of three models t also in M . We shall prove that any 2-compatible model with M is in M . We prove this inductively, by showing that any 2-compatible model is in fact n-compatible. Consider a model m k-compatible with M and a 1)-tuple of positions in this model. The k distinct k-tuples of this (k + 1)-tuple agree with some model in M . Take three of those not necessarily distinct k models. (If the models are less than three, then .) Notice that any one of those differs in at most one position of the 1)-tuple with m. Therefore, the (k + 1)-tuple of m agrees with the majority model of those three models. Hence, m is 1)-compatible with M . Therefore, any 2-compatible model with M is in M and, by Lemma 1, M is a 2CNF set. Affine Sets. M is affine iff for any three models t is also in M . Proof. This fact follows from linear algebra and especially the theory of diophantine linear equations. It states the intuitive observation (and its converse) that every convex polytope is the convex hull of its vertices. For more on that see the book of Schrijver [Sc86]. Acknowledgments . We are grateful to Christos Papadimitriou for helpful discussions and suggestions. We are also indebted to the anonymous referees for their detailed comments and suggestions that decisively helped us improve the presentation by making it more complete and precise. --R Learning conjunctions of Horn clauses Semantical and computational considerations in Horn approximations Complexity of generalized satisfiability counting prob- lems Model Theory The complexity of theorem-proving procedures Structure identification in relational data Computers and Intractability Logical Foundations of Artificial Intelligence incremental recompilation of knowl- edge On generating all maximal independent sets Horn approximations of empirical data Reasoning with characteristic models Making believers out of computers Computational Complexity A Logic for default reasoning The complexity of satisfiability problems Theory of Linear and Integer Programming Model preference default theories Knowledge compilation using Horn approximation --TR --CTR Jean-Jacques Hbrard , Bruno Zanuttini, An efficient algorithm for Horn description, Information Processing Letters, v.88 n.4, p.177-182, November Lane A. Hemaspaandra, SIGACT news complexity theory column 34, ACM SIGACT News, v.32 n.4, December 2001 Lefteris M. Kirousis , Phokion G. Kolaitis, The complexity of minimal satisfiability problems, Information and Computation, v.187 n.1, p.20-39, November 25, Lane A. Hemaspaandra, SIGACT news complexity theory column 43, ACM SIGACT News, v.35 n.1, March 2004

coNP-completeness;boolean satisfiability;model;computational complexity;polynomial-time algorithms

298507

On Syntactic versus Computational Views of Approximability.

We attempt to reconcile the two distinct views of approximation classes: syntactic and computational. Syntactic classes such as MAX SNP permit structural results and have natural complete problems, while computational classes such as APX allow us to work with classes of problems whose approximability is well understood. Our results provide a syntactic characterization of computational classes and give a computational framework for syntactic classes.We compare the syntactically defined class MAX SNP with the computationally defined class APX and show that every problem in APX can be "placed" (i.e., has approximation-preserving reduction to a problem) in MAX SNP. Our methods introduce a simple, yet general, technique for creating approximation-preserving reductions which shows that any "well"-approximable problem can be reduced in an approximation-preserving manner to a problem which is hard to approximate to corresponding factors. The reduction then follows easily from the recent nonapproximability results for MAX SNP-hard problems. We demonstrate the generality of this technique by applying it to other classes such as MAX SNP-RMAX(2) and MIN F$^{+}\Pi_2(1)$ which have the clique problem and the set cover problem, respectively, as complete problems.The syntactic nature of MAX SNP was used by Papadimitriou and Yannakakis [J. Comput. System Sci., 43 (1991), pp. 425--440] to provide approximation algorithms for every problem in the class. We provide an alternate approach to demonstrating this result using the syntactic nature of MAX SNP. We develop a general paradigm, nonoblivious local search, useful for developing simple yet efficient approximation algorithms. We show that such algorithms can find good approximations for all MAX SNP problems, yielding approximation ratios comparable to the best known for a variety of specific MAX SNP-hard problems. Nonoblivious local search provably outperforms standard local search in both the degree of approximation achieved and the efficiency of the resulting algorithms.

Introduction The approximability of NP optimization (NPO) problems has been investigated in the past via the definition of two different types of problem classes: syntactically-defined classes such as MAX SNP, and computationally-defined classes such as APX (the class of optimization problems to which a constant factor approximation can be found in polynomial time). The former is useful for obtaining structural results and has natural complete problems, while the latter allows us to work with classes of problems whose approximability is completely determined. We attempt to develop linkages between these two views of approximation problems and thereby obtain new insights about both types of classes. We show that a natural generalization of MAX SNP renders it identical to the class APX. This is an unexpected validation of Papadimitriou and Yannakakis's definition of MAX SNP as an attempt at providing a structural basis to the study of approximability. As a side-effect, we resolve the open problem of identifying complete problems for MAX NP. Our techniques extend to a generic theorem that can be used to create an approximation hierarchy. We also develop a generic algorithmic paradigm which is guaranteed to provide good approximations for MAX SNP problems, and may also have other applications. 1.1 Background and Motivation A wide variety of classes are defined based directly on the polynomial-time approximability of the problems contained within, e.g., APX (constant-factor approximable problems), PTAS (problems with polynomial-time approximation schemes), and FPTAS (problems with fully-polynomial-time approximation schemes). The advantage of working with classes defined using approximability as the criterion is that it allows us to work with problems whose approximability is well-understood. Crescenzi and Panconesi [9] have recently also been able to exhibit complete problems for such classes, particularly APX. Unfortunately such complete problems seem to be rare and artificial, and do not seem to provide insight into the more natural problems in the class. Research in this direction has to find approximation-preserving reductions from the known complete but artificial problems in such classes to the natural problems therein, with a view to understanding the approximability of the latter. The second family of classes of NPO problems that have been studied are those defined via syntactic considerations, based on a syntactic characterization of NP due to Fagin [10]. Research in this direction, initiated by Papadimitriou and Yannakakis [22], and followed by Panconesi and Ranjan [21] and Kolaitis and Thakur [19], has led to the identification of approximation classes such as MAX SNP, RMAX(2), and (1). The syntactic prescription in the definition of these classes has proved very useful in the establishment of complete problems. Moreover, the recent results of Arora, Lund, Motwani, Sudan, and Szegedy [3] have established the hardness of approximating complete problems for MAX SNP to within (specific) constant factors unless It is natural to wonder why the hardest problems in this syntactic sub-class of APX should bear any relation to all of NP. Though the computational view allows us to precisely classify the problems based on their approxima- bility, it does not yield structural insights into natural questions such as: Why certain problems are easier to approximate than some others? What is the canonical structure of the hardest representative problems of a given approximation class? and, so on. Furthermore, intuitively speaking, this view is too abstract to identification of, and reductions to establish, natural complete problems for a class. The syntactic view, on the other hand, is essentially a structural view. The syntactic prescription gives a natural way of identifying canonical hard problems in the class and performing approximation-preserving reductions to establish complete problems. Attempts at trying to find a class with both the above mentioned properties, i.e., natural complete problems and capturing all problems of a specified approximability, have not been very successful. Typically the focus has been to relax the syntactic criteria to allow for a wider class of problems to be included in the class. However in all such cases it seems inevitable that these classes cannot be expressive enough to encompass all problems with a given approximability. This is because each of these syntactically defined approximation classes is strictly contained in the class NPO; the strict containment can be shown by syntactic considerations alone. As a result if we could show that any of these classes contains all of P, then we would have separated P from NP. We would expect that every class of this nature would be missing some problems from P, and this has indeed been the case with all current definitions. We explore a different direction by studying the structure of the syntactically defined classes when we look at their closure under approximation-preserving reductions. The advantage of this is that the closure maintains the complete problems of the set, while managing to include all of P into the closure (for problems in P, the reduction is to simply use a polynomial time algorithm to compute an exact solution). It now becomes interesting, for example, to compare the closure 1 of MAX SNP (denoted MAX SNP) with APX. A positive resolution, i.e., MAX would immediately imply the non-existence of a PTAS for MAX SNP-hard problems, since it is known that PTAS is a strict subset of APX, if P 6= NP. On the other hand, an unconditional negative result would be difficult to obtain, since it would imply P 6= NP. Here we resolve this question in the affirmative. The exact nature of the result obtained depends upon the precise notion of an approximation preserving reduction used to define the closure of the class MAX SNP. The strictest notion of such reductions available in the literature are the L-reductions due to Papadimitriou and Yannakakis [22]. We introduce a new notion of reductions, called E-reductions, which are a slight extension of L-reductions. Using such reductions to define the class MAX SNP we show that this equals APX-PB, the class of all polynomially bounded NP optimization problems which are approximable to within constant factors. By using slightly looser definitions of approximation preserving reductions (and in particular the PTAS-reductions of Crescenzi et al [8]) this can be extended to include all of APX into MAX SNP. We then build upon this result to identify an interesting hierarchy of such approximability classes. An interesting side-effect of our results is the positive answer to the question of Papadimitriou and Yannakakis [22] about whether MAX NP has any complete problems. The syntactic view seems useful not only in obtaining structural complexity results but also in developing paradigms for designing efficient approximation algorithms. Exploiting the syntactic nature of MAX SNP, we develop a general paradigm for designing good approximation algorithms for problems in that class and thereby provide a more computational view of it. We refer to this paradigm as non-oblivious local search, and it is a modification of the standard local search technique [24]. We show that every MAX SNP problem can be approximated to within constant factors by such algorithms. It turns out that the performance of non-oblivious local search is comparable to that of the best-known approximation algorithms for several interesting and representative problems in MAX SNP. An intriguing possibility is that this is not a coincidence, but rather a hint at the universality of the paradigm or some variant thereof. Our results are related to some extent to those of Ausiello and Protasi [4]. They define a class GLO (for Guaranteed Local Optima) of NPO problems which have the property that for all locally optimum solutions, the ratio between the value of the global and the local optimum is bounded by a constant. It follows that GLO is a subset of APX, and it was shown that it is in fact a strict subset. We show that a MAX SNP problem is not contained in GLO, thereby establishing that MAX SNP is not contained in GLO. This contrasts with our notion of non-oblivious local search which is guaranteed to provide constant factor approximations for all problems in MAX SNP. In fact, our results indicate that non-oblivious local search is significantly more powerful than standard local search in that it delivers strictly better constant ratios, and also will provide constant factor approximations to problems not in GLO. Independently of our work, Alimonti [1] has used a similar local search technique for the approximation of a specific problem not contained in GLO or MAX SNP. 1 Papadimitriou and Yannakakis [22] hinted at the definition of MAX SNP by stating that: minimization problems will be "placed" in the classes through L-reductions to maximization problems. 1.2 Summary of Results In Section 2, we present the definitions required to state our results, and in particular the definitions of an E- reduction, APX, APX-PB, MAX SNP and MAX SNP. In Section 3, we show that MAX A generic theorem which allows to equate the closure of syntactic classes to appropriate computational classes is outlined in Section 4; we also develop an approximation hierarchy based on this result. The notion of non-oblivious local search and NON-OBLIVIOUS GLO is developed in Section 5. In Section 6, we illustrate the power of non-obliviousness by first showing that oblivious local search can achieve at most the performance ratio 3=2 for MAX 2-SAT, even if it is allowed to search exponentially large neighborhoods; in contrast, a very simple non-oblivious local search algorithm achieves a performance ratio of 4=3. We then establish that this paradigm yields a 2 k approximation to MAX k-SAT. In Section 7, we provide an alternate characterization of MAX SNP via a class of problems called MAX k-CSP. It is shown that a simple non-oblivious algorithm achieves the best-known approximation for this problem, thereby providing a uniform approximation for all of MAX SNP. In Section 8, we further illustrate the power of this class of algorithms by showing that it can achieve the best-known ratio for a specific MAX SNP problem and for VERTEX COVER (which is not contained in GLO). This implies that MAX SNP is not contained in GLO, and that GLO is strict subset of NON-OBLIVIOUS GLO. In Section 9, we apply it to approximating the traveling salesman problem. Finally, in Section 10, we apply this technique to improving a long-standing approximation bound for maximum independent sets in bounded-degree graphs. A preliminary version of this paper appeared in [18]. Preliminaries and Definitions Given an NPO problem P and an instance I of P, we use jIj to denote the length of I and OPT (I) to denote the optimum value for this instance. For any solution S to I, the value of the solution, denoted by V (I; S), is assumed to be a polynomial time computable function which takes positive integer values (see [7] for a precise definition of NPO). solution S to an instance I of an NPO problem P has error E(I; Notice that the above definition of error applies uniformly to the minimization and maximization problems at all levels of approximability. (Performance Ratio) An approximation algorithm A for an optimization problem P has performance ratio R(n) if, given an instance I of P with oe A solution of value within a multiplicative factor r of the optimal value is referred to as an r-approximation. The performance ratio for A is R if it always computes a solution with error at most R \Gamma 1. 2.1 E-reductions We now describe the precise approximation preserving reduction we will use in this paper. Various other notions of approximation preserving reductions exist in the literature (cf. [2, 16]) but the reduction which we use, referred to as the E-reduction (for error-preserving reduction), seems to be the strictest. As we will see, the E-reduction is essentially the same as the L-reduction of Papadimitriou and Yannakakis [22] and differs from it in only one relatively minor aspect. problem P E-reduces to a problem P 0 (denoted exist polynomial time computable functions f , g and a constant fi such that maps an instance I of P to an instance I 0 of P 0 such that OPT (I) and OPT are related by a polynomial factor, maps solutions S 0 of I 0 to solutions S of I such that E(I; Remark 1 An E-reduction is essentially the strictest possible notion of reduction. It requires that the error for P be linearly related to the error for P 0 . Most other notions of reductions in the literature, for example the F -reductions and P -reductions of Crescenzi and Panconesi [9], do not enforce this condition. One important consequence of this constraint is that E-reductions are sensitive, i.e., when I 2 P is mapped to I under an E-reduction then a good solution to I 0 should provide structural information about a good solution to I. Thus, reductions from real optimization problems to decision problems artificially encoded as optimization problems are implausible. Having P /E P 0 implies that P is as well approximable as P 0 ; in fact, an E-reduction is an FPTAS-preserving reduction. An important benefit is that this reduction can applied uniformly at all levels of approximability. This is not the case with the other existing definitions of FPTAS-preserving reduction in the literature. For example, the FPTAS-preserving reduction of Crescenzi and Panconesi [9] is much more unrestricted in scope and does not share this important property of the E-reduction. Note that Crescenzi and Panconesi [9] showed that there exists a problem P 0 2 PTAS such that for any problem . Thus, there is the undesirable situation that a problem P with no PTAS has a FPTAS-preserving reduction to a problem P 0 with a PTAS. Remark 3 The L-reduction of Papadimitriou and Yannakakis [22] enforces the condition that the optima of an instance I of P be linearly related to the optima of the instance I 0 of P 0 to which it is mapped. This appears to be an unnatural restriction considering that the reduction itself is allowed to be an arbitrary polynomial time computation. This is the only real difference between their L-reduction and our E-reduction, and an E-reduction in which the linearity relation of the optimas is satisfied is an L-reduction. Intuitively, however, in the study of approximability the desirable attribute is simply that the errors in the corresponding solutions are closely (linearly) related. The somewhat artificial requirement of a linear relation between the optimum values precludes reductions between problems which are related to each other by some scaling factor. For instance, it seems desirable that two problems whose objective functions are simply related by any fixed polynomial factor should be inter-reducible under any reasonable definition of an approximation-preserving reduction. Our relaxation of the L-reduction constraint is motivated precisely by this consideration. Let C be any class of NPO problems. Using the notion of an E-reduction, we define hardness and completeness of problems with respect C, as well its closure and polynomially-bounded sub-class. Definition 4 (Hard and Complete Problems) A problem P 0 is said to be C-hard if for all problems we have P /E P 0 . A C-hard problem P is said to be C-complete if in addition P 2 C. Definition 5 (Closure) The closure of C, denoted by C, is the set of all NPO problems P such that P /E P 0 for some P 0 2 C. Remark 4 The closure operation maintains the set of complete problems for a class. Definition 6 (Polynomially Bounded Subset) The polynomially bounded subset of C, denoted C-PB, is the set of all problems P 2 C for which there exists a polynomial p(n) such that for all instances I 2 P, 2.2 Computational and Syntactic Classes We first define the basic computational class APX. Definition 7 (APX) An NPO problem P is in the class APX if there exists a polynomial time computable function A mapping instances of P to solutions, and a constant c - 1, such that for all instances I of P, The class APX-PB consists of all polynomially bounded NPO problems which can be approximated within constant factors in polynomial time. If we let F-APX denote the class of NPO problems that are approximable to within a factor F , then we obtain a hierarchy of approximation classes. For instance, poly-APX and log-APX are the classes of NPO problems which have polynomial time algorithms with performance ratio bounded polynomially and logarithmically, respectively, in the input length. A more precise form of these definitions are provided in Section 4. Let us briefly review the definition of some syntactic classes. Definition 8 (MAX SNP and MAX NP [22]) MAX SNP is the class of NPO problems expressible as finding the structure S which maximizes the objective function denotes the input (consisting of a finite universe U and a finite set of bounded arity predicates P ), S is a finite structure, and F is a quantifier-free first-order formula. The class MAX NP is defined analogously except the objective function is A natural extension is to associate a weight with every tuple in the range of the universal quantifier; the modified objective is to find an S which maximizes V (I; ~x w(~x)F(I; the weight associated with the tuple ~x. Example 1 (MAX k-SAT) The MAX k-SAT problem is: given a collection of m clauses on n boolean variables where each (possibly weighted) clause is a disjunction of precisely k literals, find a truth assignment satisfying a maximum weight collection of clauses. For any fixed integer k, MAX k-SAT belongs to the class MAX SNP. The results of Papadimitriou and Yannakakis [22] can be adapted to show that for k - 2, MAX k-SAT is complete under E-reductions for the class MAX SNP. Definition 9 (RMAX(k) [21]) RMAX(k) is the class of NPO problems expressible as finding a structure which maximizes the objective function where S is a single predicate and F(I; S; ~y) is a quantifier-free CNF formula in which S occurs at most k times in each clause and all its occurrences are negative. The results of Panconesi and Ranjan [21] can be adapted to show that MAX CLIQUE is complete under E-reductions for the class RMAX(2). is the class of NPO problems expressible as finding a structure S which minimizes the objective function where S is a single predicate, F(I; S; ~y) is a quantifier-free CNF formula in which S occurs at most k times in each clause and all its occurrences are positive. The results of Kolaitis and Thakur [19] can be adapted to show that SET COVER is complete under E-reductions for the class MIN F 3 MAX SNP Closure and APX-PB In this section, we will establish the following theorem and examine its implications. The proof is based on the results of Arora et al [3] on efficient proof verifications. Theorem 1 MAX Remark 5 The seeming weakness that MAX SNP only captures polynomially bounded APX problems can be removed by using looser forms of approximation-preserving reduction in defining the closure. In particular, Crescenzi and Trevisan [8] define the notion of a PTAS-preserving reduction under which APX-PB. Using their result in conjunction with the above theorem, it is easily seen that MAX This weaker reduction is necessary to allow for reductions from fine-grained optimization problems to coarser (polynomially-bounded) optimization problems (cf. [8]). The following is a surprising consequence of Theorem 1. Theorem 2 MAX Papadimitriou and Yannakakis [22] (implicitly) introduced both these closure classes but did not conjecture them to be the same. It would be interesting to see if this equality can be shown independent of the result of Arora et al [3]. We also obtain the following resolution to the problem posed by Papadimitriou and Yannakakis [22] of finding complete problems for MAX NP. Theorem 3 MAX SAT is complete for MAX NP. The following sub-sections establish that MAX SNP ' APX-PB. The idea is to first E-reduce any minimization problem in APX-PB to a maximization problem in therein, and then E-reduce any maximization problem in APX-PB to a specific complete problem for MAX SNP, viz., MAX 3-SAT. Since an E-reduction forces the optimas of the two problems involved to be related by polynomial factors, it is easy to see that MAX SNP ' APX-PB. Combining, we establish Theorem 1. 3.1 Reducing Minimization to Maximization Observe that the fact that P belongs to APX implies the existence of an approximation algorithm A and a constant c such that c Henceforth, we will use a(I) to denote V (I; A(I)). We first reduce any minimization problem P 2 APX-PB to a maximization problem P 0 2 APX-PB, where the latter is obtained by merely modifying the objective function for P, as follows: let P 0 have the objective function V 0 (I; instances I and solutions S for P. It can be verified that the optimum value for any instance I of P 0 always lies between a(I) and (c 1)a(I). Thus, A is a 1)-approximation algorithm for P 0 . If S is a ffi -error solution to the optimum of P 0 , i.e., where OPT 0 (I) is the optimal value of V 0 for I. We obtain that c c c Thus a solution s to P 0 with error ffi is a solution to P with error at most (c implying an E-reduction with 3.2 NP Languages and MAX 3-SAT The following theorem, adapted from a result of Arora, Lund, Motwani, Sudan, and Vazirani [3], is critical to our E-reduction of maximization problems to MAX 3-SAT. Theorem 4 Given a language L 2 NP and an instance x 2 S n , one can compute in polynomial time an instance F x of MAX 3-SAT, with the following properties. 1. The formula F x has m clauses, where m depends only on n. 2. There exists a constant ffl ? 0 such that (1 \Gamma ffl)m clauses of F x are satisfied by some truth assignment. 3. If x 2 L, then F x is (completely) satisfiable. 4. If x 62 L, then no truth assignment satisfies more than clauses of F x . 5. Given a truth assignment which satisfies more than clauses of F x , a truth assignment which satisfies F x completely can be constructed in polynomial time. Some of the properties above may not be immediately obvious from the construction given by Arora, Lund, Motwani, Sudan, and Szegedy [3]. It is easy to verify that they provide a reduction with properties (1), (3) and (4). Property (5) is obtained from the fact that all assignments which satisfy most clauses are actually close (in terms of Hamming distance) to valid codewords from a linear code, and the uniquely error-corrected codeword obtained from this "corrupted code-word" will satisfy all the clauses of F x . Property (2) requires a bit more care and we provide a brief sketch of how it may be ensured. The idea is to revert back to the PCP model and redefine the proof verification game. Suppose that the original game had the properties that for x 2 L there exists a proof such that the verifier accepts with probability 1, and otherwise, for x 62 L, the verifier accepts with probability at most 1=2. We now augment this game by adding to the proof a 0th bit which the prover uses as follows: if the bit is set to 1, then the prover "chooses" to play the old game, else he is effectively "giving up" on the game. The verifier in turn first looks at the 0th bit of the proof. If this is set, then she performs the usual verification, else she tosses an unbiased coin and accepts if and only if it turns up heads. It is clear that for x 2 L there exists a proof on which the verifier always accepts. Also, for x 62 L no proof can cause the verifier to accept with probability greater than 1=2. Finally, by setting the 0th bit to 0, the prover can create a proof which the verifier accepts with probability exactly 1=2. This proof system can now be transformed into a 3-CNF formula of the desired form. 3.3 Reducing Maximization to MAX 3-SAT We have already established that, without loss of generality, we only need to worry about maximization problems Consider such a problem P, and let A be a polynomial-time algorithm which delivers a c-approximation for P, where c is some constant. Given any instance I of P, let the bound on the optimum value for I obtained by running A on input I. Note that this may be a stronger bound than the a priori polynomial bound on the optimum value for any instance of length jIj. An important consequence is that p - c OPT (I). We generate a sequence of NP decision problems L Given an instance I, we create p formulas F i , for using the reduction from Theorem 4, where ith formula is obtained from the NP language L i . Consider now the formula that has the following features. ffl The number of satisfiable clauses of F is exactly where ffl and m are as guaranteed by Theorem 4. ffl Given an assignment which satisfies (1 clauses of F , we can construct in polynomial time a solution to I of value at least j. To see this, observe the following: any assignment which so many clauses must satisfy more than clauses in at least j of the formulas F i . Let i be the largest index for which this happens; clearly, i - j. Furthermore, by property (5) of Theorem 4, we can now construct a truth assignment which satisfies F i completely. This truth assignment can be used to obtain a solution S such that V (I; In order to complete the proof it remains to be shown that given any truth assignment with error ffi , i.e., which satisfies MAX =(1 clauses of F , we can find a solution S for I with error E(I; some constant fi. We show that this is possible for cffl)=ffl. The main idea behind finding such a solution is to use the second property above to find a "good" solution to I using a "good" truth assignment for F . Suppose we are given a solution which satisfies MAX =(1 clauses. Since MAX =(1 we can use the second feature from above to construct a solution S 1 such that fflm c readily seen that Assuming we obtain that On the other hand, if then the error in a solution S 2 obtained by running the c-approximation algorithm for P is given by Therefore, choosing immediately obtain that the solution with larger value, among S 1 has error at most fi ffi . Thus, this reduction is indeed an E-reduction. 4 Generic Reductions and an Approximation Hierarchy In this section we describe a generic technique for turning a hardness result into an approximation preserving reduction. We start by listing the kind of constraints imposed on the hardness reduction, the approximation class and the optimization problem. We will observe at the end that these restrictions are obeyed by all known hardness results and the corresponding approximation classes. Definition 11 (Additive Problems) An NPO problem P is said to be additive if there exists an operator which maps a pair of instances I 1 and I 2 to an instance I 1 I 2 such that OPT Definition 12 (Downward Closed Family) A family of functions is said to be downward closed if for all g 2 F and for all constants c, g 0 (n) 2 O(g(n c )) implies that g 0 2 F . A function g is said to be hard for the family F if for all g 0 2 F , there exists a constant c such that g 0 (n) 2 O(g(n c )); the function g is said to be complete for F if g is hard for F and g 2 F . For a downward closed family F , the class F -APX consists of all problems approximable to within a ratio of g(jIj) for some function g 2 F . Definition 14 (Canonical Hardness) An NP maximization problem P is said to be canonically hard for the class F -APX if there exists a transformation T , constants n 0 and c, and a gap function G which is hard for the family F , such that given an instance x of 3-SAT on n - n 0 variables and N - n c , I = T (x; N) is an instance of P with the following properties. ffl If x 2 3-SAT, then OPT ffl If x 62 3-SAT, then OPT ffl Given a solution S to I with V (I; S) ? N=G(N), a truth assignment satisfying x can be found in polynomial time. Canonical hardness for NP minimization problems is analogously defined: OPT when the formula is satisfiable and OPT Given any solution with value less than NG(N), one can construct a satisfying assignment in polynomial time. 4.1 The Reduction Theorem 5 If F is a downward closed family of functions, and an additive NPO problem W is canonically hard for the class F -APX, then all problems in F -APX E-reduce to P. Proof: Let P be a problem in F-APX, approximable to within c(:), and let W be a problem shown to be hard to within a factor G(:) where G is complete for F . We start with the special case where both P and W are maximization problems. We describe the functions f , g and the constant fi as required for an E-reduction. Let I 2 P be an instance of size n; pick N so that c(n) is O(G(N)). To describe our reduction, we need to specify the functions f and g. The function f is defined as follows. Let each denote the NP-language fIj OPT (I) - ig. Now for each i, we create an instance OE i 2 W of size N such that if I 2 L i then OPT (OE i ) is N , and it is N=G(N) otherwise. We define We now construct the function g. Given an instance I 2 P and a solution s 0 to f(I), we compute a solution s to I in the following manner. We first use A to find a solution s 1 . We also compute a second solution s 2 to I as follows. Let j be the largest index such that the solution s 0 projects down to a solution s to the instance OE j such that V 0 This in turn implies we can find a solution s 2 to witness Our solution s is the one among s 1 and s 2 that yields the larger objective function value. We now show that the reduction holds for c(n) Consider the following two cases: Case 1 [j - m]: In this case, V (I; m. Thus s is an (ff \Gamma 1) approximate solution to I. We now argue that s 0 is at best a (ff \Gamma 1)=fi approximate solution to OE. We start with the following upper bound on c(n) G(N) \GammaG(N) Thus the approximation factor achieved by s 0 is given by Nm So in this case s 1 (and hence s) approximates I to within a factor of fi ffl, if s 0 approximates OE to within a factor of ffl. Case 2 [j - m]: Let flm. Note that fl ? 1 and that s is an (ff \Gamma fl)=fl approximate solution to I. We bound the value of the solution s 0 to OE as c(n) and its quality as Thus in this case also we find that s (by virtue of s 2 ) is a solution of quality fi ffl if s 0 is a solution of quality s. We now consider the more general cases where P and W are not both maximization problems. For the case where both are minimization problems, the above transformation works with one minor change. When creating OE i , the NP language consists of instances (I; i) such that there exists s with For the case where P is a minimization problem and W is a maximization problem, we first E-reduce P to a maximization problem P 0 and then proceed as before. The reduction proceeds as follows. The objective function of P 0 is defined as V 0 (I; To begin with, it is easy to verify that P 2 F-APX implies Let s be a fi approximate solution to instance I of P. We will show that s is at best a fi=2 approximate solution to instance I of P 0 . Assume, without loss of generality, that fi 6= 0. Then Multiplying by 2m 2 =(OPT (I)V (I; s)), we get 2: This implies that Upon rearranging, Thus the reduction from P to P 0 is an E-reduction. Finally, the last remaining case, i.e., P being a maximization problem and W being a minimization problem, is dealt with similarly: we transform P into a minimization problem P 0 . Remark 6 This theorem appears to merge two different notions of the relative ease of approximation of optimization problems. One such notion would consider a problem P 1 easier than P 2 if there exists an approximation preserving reduction from P 1 to P 2 . A different notion would regard P 1 to be easier than one seems to have a better factor of approximation than the other. The above statement essentially states that these two comparisons are indeed the same. For instance, the MAX CLIQUE problem and the CHROMATIC NUMBER problem, which are both in poly-APX, are inter-reducible to each other. The above observation motivates the search for other interesting function classes f , for which the class f -APX may contain interesting optimization problems. 4.2 Applications The following is a consequence ofTheorem 5. Theorem 6 a) b) If SET COVER is canonically hard to approximate to within a factor of W(log n), then We briefly sketch the proof of this theorem. The hardness reduction for MAX SAT and CLIQUE are canonical [3, 11]. The classes APX-PB, poly-APX, log-APX are expressible as classes F-APX for downward closed function families. The problems MAX SAT, MAX CLIQUE and SET COVER are additive. Thus, we can now apply Theorem 5. Remark 7 We would like to point out that almost all known instances of hardness results seem to be shown for problems which are additive. In particular, this is true for all MAX SNP problems, MAX CLIQUE, CHROMATIC NUMBER, and SET COVER. One case where a hardness result does not seem to directly apply to an additive problem is that of LONGEST PATH [17]. However in this case, the closely related LONGEST S-T PATH problem is easily seen to be additive and the hardness result essentially stems from this problem. Lastly, the most interesting optimization problems which do not seem to be additive are problems related to GRAPH BISECTION or PARTITION, and these also happen to be notable instances where no hardness of approximation results have been achieved! 5 Local Search and MAX SNP In this section we present a formal definition of the paradigm of non-oblivious local search, and describe how it applies to a generic MAX SNP problem. Given a MAX SNP problem P, recall that the goal is to find a structure S which maximizes the objective function: V (I; ~x F(I; S; ~x). In the subsequent discussion, we view S as a k-dimensional boolean vector. 5.1 Classical Local Search We start by reviewing the standard mechanism for constructing a local search algorithm. A ffi-local algorithm A for P is based on a distance function D(S 1 which is the Hamming distance between two k-dimensional vectors. The ffi-neighborhood of a structure S is given by N(S; U is the universe. A structure S is called ffi-optimal if 8S algorithm computes a ffi-optimum by performing a series of greedy improvements to an initial structure S 0 , where each iteration moves from the current structure S i to some S For constant ffi , a ffi-local search algorithm for a polynomially-bounded NPO problem runs in polynomial time because: ffl each local change is polynomially computable, and ffl the number of iterations is polynomially bounded since the value of the objective function improves monotonically by an integral amount with each iteration, and the optimum is polynomially-bounded. 5.2 Non-Oblivious Local Search A non-oblivious local search algorithm is based on a 3-tuple hS is the initial solution structure which must be independent of the input, F(I; S) is a real-valued function referred to as the weight function, and D is a real-valued distance function which returns the distance between two structures in some appropriately chosen metric. The distance function D is computable in time polynomial in jU j. Thus as before, for constant ffi, a non-oblivious ffi -local algorithm terminates in time polynomial in the input size. The classical local search paradigm, which we call oblivious local search, makes the natural choice for the function F(I; S), and the distance function D, i.e., it chooses them to be V (I; S) and the Hamming distance. However, as we show later, this choice does not always yield a good approximation ratio. We now formalize our notion of this more general type of local search. Definition 15 (Non-Oblivious Local Search Algorithm) A non-oblivious local search algorithm is a local search algorithm whose weight function is defined to be ~x r where r is a constant, F i 's are quantifier-free first-order formulas, and the profits p i are real constants. The distance function D is an arbitrary polynomial-time computable function. A non-oblivious local search can be implemented in polynomial time in much the same way as oblivious local search. Note that the we are only considering polynomially-bounded weight functions and the profits are fixed independent of the input size. In general, the non-oblivious weight functions do not direct the search in the direction of the actual objective function. In fact, as we will see, this is exactly the reason why they are more powerful and allow for better approximations. Definition (Non-Oblivious GLO) The class of problems NON-OBLIVIOUS GLO consists of all problems which can be approximated within constant factors by a non-oblivious ffi -local search algorithm for some constant ffi . Remark 8 We make some observations about the above definition. It would be perfectly reasonable to allow weight functions which are non-linear, but we stay with the above definition for the purposes of this paper. Allowing only a constant number of predicates in the weight functions enables us to prevent the encoding of arbitrarily complicated approximation algorithms. The structure S is a k-dimensional vector, and so a convenient metric for the distance function D is the Hamming distance. This should be assumed to be the underlying metric unless otherwise specified. However, we have found that it is sometimes useful to modify this, for example by modifying the Hamming distance so that the complement of a vector is considered to be at distance 1 from it. Finally, it is sometimes convenient to assume that the local search makes the best possible move in the bounded neighborhood, rather than an arbitrary move which improves the weight function. We believe that this does not increase the power of non-oblivious local search. 6 The Power of Non-Oblivious Local Search We will show that there exists a choice of a non-oblivious weight function for MAX k-SAT such that any assignment which is 1-optimal with respect to this weight function, yields a performance ratio of 2 k with respect to the optimal. But first, we obtain tight bounds on the performance of oblivious local search for MAX 2-SAT, establishing that its performance is significantly weaker than the best-known result even when allowed to search exponentially large neighborhoods. We use the following notation: for any fixed truth assignment ~ is the set of clauses in which exactly i literals are true; and, for a set of clauses S, W (S) denotes the total weight of the clauses in S. 6.1 Oblivious Local Search for MAX 2-SAT We show a strong separation in the performance of oblivious and non-oblivious local search for MAX 2-SAT. Suppose we use a ffi-local strategy with the weight function F being the total weight of the clauses satisfied by the assignment, i.e., The following theorem shows that for any an oblivious ffi -local strategy cannot deliver a performance ratio better than 3=2. This is rather surprising given that we are willing to allow near-exponential time for the oblivious algorithm. Theorem 7 The asymptotic performance ratio for an oblivious ffi -local search algorithm for MAX 2-SAT is 3=2 for any positive This ratio is still bounded by 5=4 when ffi may take any value less than n=2. Proof: We show the existence of an input instance for MAX 2-SAT which may elicit a relatively poor performance ratio for any ffi -local algorithm provided In our construction of such an input instance, we assume that n - 2ffi + 1. The input instance comprises of a disjoint union of four sets of clauses, say defined as below: 1-i!j-n 1-i!j-n 2ffi+2-i-n i!j-n Clearly, 1). Without loss of generality, assume that the current input assignment is ~ 1). This satisfies all clauses in G 1 and G 2 . But none of the clauses in G 3 and G 4 are satisfied. If we flip the assignment of values to any k - ffi variables, it would unsatisfy precisely k(n \Gamma clauses in G 1 . This is the number of clauses in G 1 flipped variable occurs with an unflipped variable. On the other hand, flipping the assigned values of any k - ffi variables can satisfy at most k(n \Gamma clauses in G 3 as we next show. denote the set of clauses on n variables given by We claim the following. assignment of values to the n variables such that at most k - ffi variables have been assigned value false, can satisfy at most k(n \Gamma clauses in P(n; ffi). Proof: We prove by simultaneous induction on n and ffi that the statement is true for any instance P(n; ffi) where n and ffi are non-negative integers such that n. The base case includes and is trivially verified to be true for the only allowable value of ffi , namely We now assume that the statement is true for any instance P(n Consider now the instance P(n; ffi ). The statement is trivially true for g be any choice of k - ffi variables such that q. Again the assertion is trivially true if We assume that k - 2 from now on. If we delete all clauses containing the variables z 1 and z 2 from P(n; ffi ), we get the instance P(n \Gamma 2; 1). We now consider three cases. Case 1 In this case, we are reduced to the problem of finding an upper bound on the maximum number of clauses satisfied by setting any k variables to false in P(n \Gamma 2; use the inductive hypothesis to conclude that no more than (n clauses will be satisfied. Thus the assertion holds in this case. However, we may not directly use the inductive hypothesis if But in this case we observe that since by the inductive hypothesis, setting any k \Gamma 1 variables in P(n \Gamma 2; false, satisfies at most (n clauses, assigning the value false to any set of k variables, can satisfy at most clauses. Hence the assertion holds in this case also. Case In this case, z j 1 satisfies one clause and the remaining variables satisfy at most clauses by the inductive hypothesis on Adding up the two terms, we see that the assertion holds. Case 3 We analyze this case based on whether precisely clauses and the remaining variables, satisfy at most (n clauses using the inductive hypothesis. Thus the assertion still holds. Otherwise, z 1 satisfies precisely clauses and the remaining no more than (n clauses using the inductive hypothesis. Summing up the two terms, we get (n \Gamma k)k as the upper bound on the total number of clauses satisfied. Thus the assertion holds in this case also. To see that this bound is tight, simply consider the situation when the k variables set to false are ffi. The total number of clauses satisfied is given by Assuming that each clause has the same weight, Lemma 1 allows us to conclude that a ffi-local algorithm cannot increase the total weight of satisfied clauses with this starting assignment. An optimal assignment on the other hand can satisfy all the clauses by choosing the vector ~ 0). Thus the performance ratio of a ffi-local algorithm, say R ffi , is bounded as For any asymptotically converges to 3=2. We next show that this bound is tight since a 1-local algorithm achieves it. However, before we do so, we make another intriguing observation, namely, for any ffi ! n=2, the ratio R ffi is bounded by 5=4. To see that a 1-local algorithm ensures a performance ratio of 3=2, consider any 1-optimal assignment ~ Z and let ff i denote the set of clauses containing the variable z i such that no literal in any clause of ff i is satisfied by ~ Z . Similarly, let fi i denote the set of clauses containing the variable z i such that precisely one literal is satisfied in any clause in fi i and furthermore, it is precisely the literal containing the variable z i . If we complement the value assigned to the variable z i , it is exactly the set of clauses in ff i which becomes satisfied and the set of clauses in fi i which is no longer satisfied. Since ~ Z is 1-optimal, it must be the case that W (ff i ) - W (fi i ). If we sum up this inequality over all the variables, then we get the inequality We observe that because each clause in S 0 gets counted twice while each clause in S 1 gets counted exactly once. Thus the fractional weight of the number of clauses not satisfied by a 1-local assignment is bounded as Hence the performance ratio achieved by a 1-local algorithm is bounded from above by 3=2. Combining this with the upper bound derived earlier, we conclude that R We may summarize our results as follows. Lemma 2 The performance ratio R ffi for any ffi -local algorithm for MAX 2-SAT using the weight function positive integer o(n). Furthermore, this ratio is still bounded by 5=4 when ffi may take any value less than n=2. 6.2 Oblivious Local Search for MAX 2-SAT We now illustrate the power of non-oblivious local search by showing that it achieves a performance ratio of 4=3 for MAX 2-SAT, using 1-local search with a simple non-oblivious weight function. Theorem 8 Non-oblivious 1-local search achieves a performance ratio of 4=3 for MAX 2-SAT. Proof: We use the non-oblivious weight function Consider any assignment ~ Z which is 1-optimal with respect to this weight function. Without loss of generality, we assume that the variables have been renamed such that each unnegated literal gets assigned the value true. Let P i;j and N i;j respectively denote the total weight of clauses in S i containing the literals z j and z j , respectively. Since ~ Z is a 1-optimal assignment, each variable z j must satisfy the following equation. \Gamma2 P 2;j \Gamma2 P 1;j +2 N 1;j +2 N 0;j - 0: Summing this inequality over all the variables, and using we obtain the following inequality: This immediately implies that the total weight of the unsatisfied clauses at this local optimum is no more than 1=4 times the total weight of all the clauses. Thus, this algorithm ensures a performance ratio of 4=3. Remark 9 The same result can be achieved by using the oblivious weight function, and instead modifying the distance function so that it corresponds to distances in a hypercube augmented by edges between nodes whose addresses are complement of each other. 6.3 Generalization to MAX k-SAT We can also design a non-oblivious weight function for MAX k-SAT such that a 1-local strategy ensures a performance ratio of 2 k 1). The weight function F will be of the form the coefficients c i 's will be specified later. Theorem 9 Non-oblivious 1-local search achieves a performance ratio of 2 k Proof: Again, without loss of generality, we will assume that the variables have been renamed so that each unnegated literal is assigned true under the current truth assignment. Thus the set S i is the set of clauses with i unnegated literals. denote the change in the current weight when we flip the value of z j , that is, set it to 0. It is easy to verify the following equation: (D Thus when the algorithm terminates, we know that @F Summing over all values of j, and using the fact get the following inequality. We now determine the values of D i 's such that the coefficient of each term on the left hand side is unity. It can be verified that achieves this goal. Thus the coefficient of W (S 0 ) on the right hand side of equation (2) is the weight of the clauses not satisfied is bounded by 1=2 k times the total weight of all the clauses. It is worthwhile to note that this is regardless of the value chosen for the coefficient c 0 . 7 Local Search for CSP and MAX SNP We now introduce a class of constraint satisfaction problems such that the problems in MAX SNP are exactly equivalent to the problems in this class. Furthermore, every problem in this class can be approximated to within a constant factor by a non-oblivious local search algorithm. 7.1 Constraint Satisfaction Problems The connection between the syntactic description of optimization problems and their approximability through non-oblivious local search is made via a problem called MAX k-CSP which captures all the problems in MAX SNP as a special case. Definition 17 (k-ary Constraint) Let be a set of boolean variables. A k-ary constraint on Z is is a size k subset of Z, and is a k-ary boolean predicate. Definition Given a collection C of weighted k-ary constraints over the variables g, the MAX k-CSP problem is to find a truth assignment satisfying a maximum weight sub-collection of the constraints. The following theorem shows that MAX k-CSP problem is a "universal" MAX SNP problem, in that it contains as special cases all problems in MAX SNP. Theorem a) For fixed k, MAX k-CSP 2 MAX SNP. Moreover, the k-CSP instance corresponding to any instance of this problem can be computed in polynomial time. 7.2 Non-Oblivious Local Search for MAX k-CSP A suitable generalization of the non-oblivious local search algorithm for MAX k-SAT yields the following result. Theorem 11 A non-oblivious 1-local search algorithm has performance ratio 2 k for MAX k-CSP. Proof: We use an approach similar to the one used in the previous section to design a non-oblivious weight function F for the weighted version of the MAX k-CSP problem such that a 1-local algorithm yields performance ratio to this problem. We consider only the constraints with at least one satisfying assignment. Each such constraint can be replaced by a monomial which is the conjunction of some k literals such that when the monomial evaluates to true the corresponding literal assignment represents a satisfying assignment for the constraint. Furthermore, each such monomial has precisely one satisfying assignment. We assign to each monomial the weight of the constraint it represents. Thus any assignment of variables which satisfies monomials of total weight W 0 , also satisfies constraints in the original problem of total weight W 0 . denote the monomials with i true literals, and assume that the weight function F is of the form assuming that the variables have been renamed so that the current assignment gives value true to each variable, we know that for any variable z j , @F is given by equation (1). As before, using the fact that for any 1-optimal assignment, @F summing over all values of j, we can write the following inequality. We now determine the values of D i 's such that the coefficient of each term on the left hand side is unity. It can be verified that achieves this goal. Thus the coefficient of W (S k ) on the right hand side of equation (1) is the total weight of clauses satisfied is at least 1=2 k times the total weight of all the clauses with at least one satisfiable assignment. We conclude the following theorem. Theorem 12 Every optimization problem P 2 MAX SNP can be approximated to within some constant factor by a (uniform) non-oblivious 1-local search algorithm, i.e., For a problem expressible as k-CSP, the performance ratio is at most 2 k . 8 Non-Oblivious versus Oblivious GLO In this section, we show that there exist problems for which no constant factor approximation can be obtained by any ffi -local search algorithm with oblivious weight function, even when we allow ffi to grow with the input size. However, a simple 1-local search algorithm using an appropriate non-oblivious weight function can ensure a constant performance ratio. 8.1 MAX 2-CSP The first problem is an instance of MAX 2-CSP where we are given a collection of monomials such that each monomial is an "and" of precisely two literals. The objective is to find an assignment to maximize the number of monomials satisfied. We show an instance of this problem such that for every there exists an instance one of whose local optima has value that is a vanishingly small fraction of the global optimum. The input instance consists of a disjoint union of two sets of monomials, say G 1 and G 2 , defined as below: 1-i!j-n i!j-n Clearly, . Consider the truth assignment ~ 1). It satisfies all monomials in G 2 but none of the monomials in G 1 . We claim that this assignment is ffi -optimal with respect to the oblivious weight function. To see this, observe that complementing the value of any p - ffi variables will unsatisfy at least ffip=2 monomials in G 2 for any On the other hand, this will satisfy precisely . For any p - ffi , we have (ffip)=2 - so Z is a ffi -local optimum. The optimal assignment on the other hand, namely ~ all monomials in G 1 . Thus, for n=2, the performance ratio achieved by any ffi -local algorithm is no more than which asymptotically diverges to infinity for any We have already seen in Section 7 that a 1- local non-oblivious algorithm ensures a performance ratio of 4 for this problem. Since this problem is in MAX SNP, we obtain the following theorem. Theorem 13 There exist problems in MAX SNP such that for oblivious algorithm can approximate them to within a constant performance ratio, i.e., MAX SNP 6' GLO: 8.2 Vertex Cover Ausiello and Protasi [4] have shown that VERTEX COVER does not belong to the class GLO and, hence, there does not exist any constant ffi such that an oblivious ffi-local search algorithm can compute a constant factor approximation. In fact, their example can be used to show that for any the performance ratio ensured by ffi-local search asymptotically diverges to infinity. However, we show that there exists a rather simple non-oblivious weight function which ensures a factor 2 approximation via a 1-local search. In fact, the algorithm simply enforces the behavior of the standard approximation algorithm which iteratively builds a vertex cover by simply including both end-points of any currently uncovered edge. We assume that the input graph G is given as a structure (V; fEg) where V is the set of vertices and encodes the edges of the graph. Our solution is represented by a 2-ary predicate M which is iteratively constructed so as to represent a maximal matching. Clearly, the end-points of any maximal matching constitute a valid vertex cover and such a vertex cover can be at most twice as large as any other vertex cover in the graph. Thus M is an encoding of the vertex cover computed by the algorithm. The algorithm starts with M initialized to the empty relation and at each iteration, at most one new pair is included in it. The non-oblivious weight function used is as below: where Let M encode a valid matching in the graph G. 2 We make the following observations. obtained from M by either deleting an edge from it, or including an edge in which is incident on an edge of M , has the property that F(I; M 0 ) - F(I; M). Thus in a 1-local search from M , we will never move to a relation M 0 which does not encode a valid matching of G. ffl On the other hand, if a relation M 0 corresponds to the encoding of a matching in G which is larger than the matching encoded by M , then F(I; does not encode a maximal matching in G, there always exist a relation in its 1-neighborhood of larger weight than itself. These two observations, combined with the fact that we start with a valid initial matching (the empty matching), immediately allow us to conclude that any 1-optimal relation M always encodes a maximal matching in G. We have established the following. Theorem 14 A 1-local search algorithm using the above non-oblivious weight function achieves a performance ratio of 2 for the VERTEX COVER problem. Theorem 15 GLO is a strict subset of NON-OBLIVIOUS GLO: 2 It is implicit in our formulation that M will correspond to a lower triangular matrix representation of the matching edges. 9 The Traveling Salesman Problem The TSP(1,2) problem is the traveling salesman problem restricted to complete graphs where all edge weights are either 1 or 2; clearly, this satisfies the triangle inequality. Papadimitriou and Yannakakis [23] showed that this problem is hard for MAX SNP. The natural weight function for TSP(1,2), that is, the weight of the tour, can be used to show that a 4-local algorithm yields a 3=2 performance ratio. The algorithm starts with an arbitrary tour and in each iteration, it checks if there exist two disjoint edges (a; b) and (c; d) on the tour such that deleting them and replacing them with the edges (a; c) and (b; d) yields a tour of lesser cost. Theorem 4-local search algorithm using the oblivious weight function achieves a 3=2 performance ratio for TSP(1,2). Proof: Let C be a 4-optimal solution and let Let - be a permutation such that the vertices in C occur in the order v - Consider any optimal solution O. With each unit cost edge e in O, we associate a unit cost edge e 0 in C as follows. Let consider the edges e ) on C. We claim either e 1 or e 2 must be of unit cost. Suppose not, then the tour C 0 which is obtained by simply deleting both e 1 and e 2 and inserting the edges e and has cost at least one less than C. But C is 4-optimal and thus this is a contradiction. Let UO denotes the set of unit cost edges in O and let UC be the set of unit cost edges in C which form the image of UO under the above mapping. Since an edge e in UC can only be the image of unit cost edges incident on v - i in O and since O is a tour, there are at most two edges in UO which map to e 0 . Thus jU C j - jU O j=2 and hence cost(O) The above bound can be shown to be tight. Theorem 17 There exists a TSP(1,2) instance such that the optimal solution has cost n + O(1) and there exists a certain 4-optimal solution for it with cost 3n=2 +O(1). Maximum Independent Sets in Bounded Degree Graphs The input instance to the maximum independent set problem in bounded degree graphs, denoted MIS-B, is a graph G such that the degree of any vertex in G is bounded by a constant D. We present an algorithm with performance ratio ( 1)=2 for this problem when D - 10. Our algorithm uses two local search algorithms such that the larger of the two independent sets computed by these algorithms, gives us the above claimed performance ratio. We refer to these two algorithms as A 1 and A 2 . In our framework, the algorithm A 1 can be characterized as a 3-local algorithm with the weight function simply being jI Thus if we start with I initialized to empty set, it is easy to see that at each iteration, I will correspond to an independent set in G. A convenient way of looking at this algorithm is as follows. We define an swap to be the process of deleting i vertices from S and including j vertices from the set V \Gamma S to the set S. In each iteration, the algorithm A 1 performs either a 0 however, can be interpreted as j applications of 0 swaps. Thus the algorithm may be viewed as executing a 0 swap at each iteration. The algorithm terminates when neither of these two operations is applicable. Let I denote the 3-optimal independent set produced by the algorithm A 1 . Furthermore, let O be any optimal independent set and let We make the following useful observations. ffl Since for no vertex in I , a 0 can be performed, it implies that each vertex in V \Gamma I must have at least one incoming edge to I . ffl Similarly, since no 1 swaps can be performed, it implies that at most jI \Gamma X j vertices in O \Gamma I can have precisely one edge coming into I . Thus vertices in O \Gamma X must have at least two edges entering the set I . A rather straightforward consequence of these two observations is the following lemma. Lemma 3 The algorithm A 1 has performance ratio (D + 1)=2 for MIS-B. Proof: The above two observations imply that the minimum number of edges entering I from the vertices in O \Gamma X is On the other hand, the maximum number of edges coming out of the vertices in I to the vertices in O \Gamma X is bounded by jI \Gamma X jD. Thus we must have Rearranging, we get which yields the desired result. This nearly matches the approximation ratio of D=2 due to Hochbaum [15]. It should be noted that the above result holds for a broader class of graphs, viz., k-claw free graphs. A graph is called there does not exist an independent set of size k or larger such that all the vertices in the independent set are adjacent to the same vertex. Lemma 3 applies to (D Our next objective is to further improve this ratio by using the algorithm A 1 in combination with the algorithm A 2 . The following lemma uses a slightly different counting argument to give an alternative bound on the approximation ratio of the algorithm A 1 when there is a constraint on the size of the optimal solution. Lemma 4 For any real number c ! D, the algorithm A 1 has performance ratio (D \Gamma c)=2 for MIS-B when the optimal value itself is no more than Proof: As noted earlier, each vertex in V \Gamma I must have at least one edge coming into the set I and at least vertices in O must have at least two edges coming into I . Therefore, the following inequality must be satisfied: Thus Finally, observe that The above lemma shows that the algorithm A 1 yields a better approximation ratio when the size of the optimal independent set is relatively small. The algorithm A 2 is simply the classical greedy algorithm. This algorithm can be conveniently included in our framework if we use directed local search. If we let N(I) denote the set of neighbors of the vertices in I , then the weight function is simply jI j(D 1). It is not difficult to see that starting with an empty independent set, a 1-local algorithm with directed search on above weight function simply simulates a greedy algorithm. The greedy algorithm exploits the situation when the optimal independent set is relatively large in size. It does so by using the fact that the existence of a large independent set in G ensures a large subset of vertices in G with relatively small average degree. The following two lemmas characterize the performance of the greedy algorithm. Lemma 5 Suppose there exists an independent set X ' V such that the average degree of vertices in X is bounded by ff. Then for any ff - 1, the greedy algorithm produces an independent set of size at least Proof: The greedy algorithm iteratively chooses a vertex of smallest degree in the remaining graph and then deletes this vertex and all its neighbors from the graph. We examine the behavior of the greedy by considering two types of iterations. First consider the iterations in which it picks a vertex outside X . Suppose in the ith such iteration, it picks a vertex in exactly k i neighbors in the set X in the remaining graph. Since each one of these k i vertices must also have at least k i edges incident on them, we loose at least k 2 edges incident on X . Suppose only p such iterations occur and let observe that Secondly, we consider the iterations when the greedy selects a vertex in X . Then we do not loose any other vertices in X because X is an independent set. Thus the total size of the independent set constructed by the greedy algorithm is at least p By the Cauchy-Schwartz inequality, Therefore, we have (1 Rearranging, we obtain that Thus and the result follows. Lemma 6 For non-negative real number c - 3D \Gamma has performance ratio (D \Gamma c)=2 for MIS-B when the optimal value itself is at least ((D \Gamma c)jV j)=(D+c+4). Proof: Observe that the average degree of vertices in O is bounded by (jV \Gamma OjD=jOj) and thus using the fact that jOj - (D \Gamma c)jV we know that the algorithm A 2 computes an independent set of size at least jOj=(1 Hence it is sufficient to determine the range of values c can take such that the following inequality is satisfied: Substituting the bound on the value of ff and rearranging the terms of the equation, yields the following Since c must be strictly bounded by D, the above quadratic equation is satisfied for any choice of Combining the results of Lemmas 4 and 6 and choosing the largest allowable value for c, we get the following result. Theorem approximation algorithm which simply outputs the larger of the two independent sets computed by the algorithms A 1 and A 2 , has performance ratio ( The performance ratio claimed above is essentially D=2:414. This improves upon the long-standing approximation ratio of D=2 due to Hochbaum [15], when D - 10. However, very recently, there has been a flurry of new results for this problem. Berman and Furer [6] have given an algorithm with performance (D when D is even, and (D fixed constant. Halldorsson and Radhakrishnan [14] have shown that algorithm A 1 when run on k-clique free graphs, yields an independent set of size at least 2n=(D They combine this algorithm with a clique-removal based scheme to achieve a performance ratio of D=6(1 Acknowledgements Many thanks to Phokion Kolaitis for his helpful comments and suggestions. Thanks also to Giorgio Ausiello and Pierluigi Crescenzi for guiding us through the intricacies of approximation preserving reductions and the available literature on it. --R New Local Search Approximation Techniques for Maximum Generalized Satisfiability Problems. Approximate Solution of NP Optimization Problems. Proof Verification and Hardness of Approximation Problems. Optimization Problems and Local Optima. Efficient probabilistically checkable proofs. Approximating Maximum Independent Set in Bounded Degree Graphs. Introduction to the Theory of Complexity. Completeness in approximation classes. Generalized First-Order Spectra and Polynomial-time Recognizable Sets Computers and Intractability: A Guide to the Theory of NP- Completeness Improved Approximations of Independent Sets in Bounded-Degree Graphs Efficient bounds for the stable set On the Approximability of NP-complete Optimization Problems On approximating the longest path in a graph. On Syntactic versus Computational Views of Approx- imability Approximation Properties of NP Minimization Classes. On the hardness of approximating minimization problems. Quantifiers and Approximation. The traveling salesman problem with distances one and two. The analysis of local search problems and their heuristics. --TR --CTR Bruno Escoffier , Vangelis Th. Paschos, Completeness in approximation classes beyond APX, Theoretical Computer Science, v.359 n.1, p.369-377, 14 August 2006 Angel , Evripidis Bampis , Laurent Gourvs, Approximating the Pareto curve with local search for the bicriteria TSP(1,2) problem, Theoretical Computer Science, v.310 n.1-3, p.135-146, 01 January 2004 James B. Orlin , Abraham P. Punnen , Andreas S. Schulz, Approximate local search in combinatorial optimization, Proceedings of the fifteenth annual ACM-SIAM symposium on Discrete algorithms, January 11-14, 2004, New Orleans, Louisiana Jukka Suomela, Approximability of identifying codes and locating--dominating codes, Information Processing Letters, v.103 n.1, p.28-33, June, 2007 Cristina Bazgan , Bruno Escoffier , Vangelis Th. Paschos, Completeness in standard and differential approximation classes: poly-(D)APX- and (D)PTAS-completeness, Theoretical Computer Science, v.339 n.2, p.272-292, 12 June 2005 Friedrich Eisenbrand , Fabrizio Grandoni, On the complexity of fixed parameter clique and dominating set, Theoretical Computer Science, v.326 n.1-3, p.57-67, 20 October 2004 Alex characterization of NP with optimal amortized query complexity, Proceedings of the thirty-second annual ACM symposium on Theory of computing, p.191-199, May 21-23, 2000, Portland, Oregon, United States Jianer Chen , Xiuzhen Huang , Iyad A. Kanj , Ge Xia, Polynomial time approximation schemes and parameterized complexity, Discrete Applied Mathematics, v.155 n.2, p.180-193, January, 2007 Moni Naor , Kobbi Nissim, Communication preserving protocols for secure function evaluation, Proceedings of the thirty-third annual ACM symposium on Theory of computing, p.590-599, July 2001, Hersonissos, Greece Harry B. Hunt, III , Madhav V. Marathe , Venkatesh Radhakrishnan , S. S. Ravi , Daniel J. Rosenkrantz , Richard E. Stearns, Parallel approximation schemes for a class of planar and near planar combinatorial optimization problems, Information and Computation, v.173 n.1, p.40-63, February 25, 2002

approximation algorithms;computational classes;computational complexity;local search;complete problems;polynomial reductions

298555

A Spectral Algorithm for Seriation and the Consecutive Ones Problem.

In applications ranging from DNA sequencing through archeological dating to sparse matrix reordering, a recurrent problem is the sequencing of elements in such a way that highly correlated pairs of elements are near each other. That is, given a correlation function f reflecting the desire for each pair of elements to be near each other, find all permutations $\pi$ with the property that if $\pi(i)<\pi(j)<\pi(k)$ then $f(i,j) \ge f(i,k)$ and $f(j,k) \ge f(i,k)$. This seriation problem is a generalization of the well-studied consecutive ones problem. We present a spectral algorithm for this problem that has a number of interesting features. Whereas most previous applications of spectral techniques provide only bounds or heuristics, our result is an algorithm that correctly solves a nontrivial combinatorial problem. In addition, spectral methods are being successfully applied as heuristics to a variety of sequencing problems, and our result helps explain and justify these applications.

Introduction . Many applied computational problems involve ordering a set so that closely coupled elements are placed near each other. This is the underlying problem in such diverse applications as genomic sequencing, sparse matrix envelope reduction and graph linear arrangement as well as less familiar settings such as archaeological dating. In this paper we present a spectral algorithm for this class of problems. Unlike traditional combinatorial methods, our approach uses an eigenvector of a matrix to order the elements. Our main result is that this approach correctly solves an important ordering problem we call the seriation problem which includes the well known consecutive ones problem [5] as a special case. More formally, we are given a set of n elements to sequence; that is, we wish to bijectively map the elements to the integers We also have a symmetric, real valued correlation function (sometimes called a similarity function) that reflects the desire for elements to be near each other in the sequence. We now wish to find all ways to sequence the elements so that the correlations are consistent; that is, if - is our permutation of elements k). Although there may be an exponential number of such orderings, they can all be described in a compact data structure known as a PQ-tree [5] which we review in the next section. Not all correlation functions allow for a consistent sequencing. If a consistent ordering is possible we will say the problem is well posed. Determining an ordering from a correlation function is what we will call the seriation problem, reflecting its origins in archaeology [29, 33]. This work was supported by the Mathematical, Information and Computational Sciences Division of the U.S. DOE, Office of Energy Research, and work at Sandia National Laboratories, operated for the U.S. DOE under contract No. DE-AC04-94AL85000. Sandia is multiprogram laboratory operated by Sandia Corporation, a Lockheed Martin Company, for the U.S. DOE. y Dept. Mathematics, University of Michigan, 2072 East Hall, Ann Arbor, MI 48109. [email protected]. z Scientific Computing & Computational Mathematics, Gates Bldg. 2B, Stanford Univ., Stanford, 94305-9025. [email protected]. x Applied & Numerical Mathematics Dept., Sandia National Labs, Albuquerque, NM 87185-1110. [email protected]. The consecutive ones problem (C1P) is a closely related ordering problem. A 1)-matrix C has the consecutive ones property if there exists a permutation matrix \Pi such that for each column in \PiC, all the ones form a consecutive sequence. If a matrix has the consecutive ones property, then the consecutive ones problem is to find all such permutations. As shown by Kendall [19] and reviewed in x6, C1P is a special case of the seriation problem. Our algorithm orders elements using their value in an eigenvector of a Laplacian matrix which we formally define in x2. Eigenvectors related to graphs have been studied since the 1950's (see, for example, the survey books by Cvetkovi'c et al. [8, 7]. Most of the early work involved eigenvectors of adjacency matrices. Laplacian eigenvectors were first studied by Fiedler [10, 11] and independently by Donath and Hoffman [9]. More recently, there have been a number of attempts to apply spectral graph theory to problems in combinatorial optimization. For example, spectral algorithms have been developed for graph coloring [3], graph partitioning [9, 28] and envelope reduction [4], and more examples can be found in the survey papers of Mohar [23, 24]. However, in most previous applications, these techniques have been used to provide bounds, heuristics, or in a few cases, approximation algorithms [2, 6, 14] for NP-hard problems. There are only a small number of previous results in which eigenvector techniques have been used to exactly solve combinatorial problems including finding the number of connected components of a graph [10], coloring k-partite graphs [3], and finding stable sets (independent sets) in perfect graphs [16]. This paper describes another such application. Spectral methods are closely related to the more general method of semidefinite programming, which has been applied successfully to many combinatorial problems (e.g. MAX-CUT and MAX-2SAT[14] and graph coloring[18]). See Alizadeh[1] for a survey of semidefinite programming with applications to combinatorial optimization. Our result is important for several reasons. First, it provides new insight into the well-studied consecutive ones problem. Second, some important practical problems like envelope reduction for matrices and genomic reconstruction can be thought of as variations on seriation. For example, if biological experiments were error-free, the genomic reconstruction problem would be precisely C1P. Unfortunately, real experimental data always contain errors, and attempts to generalize the consecutive ones concept to data with errors seems to invariably lead to NP-complete problems [31, 15]. A spectral heuristic based upon our approach has recently been applied to such problems and found to be highly successful in practice [15]. Our result helps explain this empirical success by revealing that in the error-free case the technique will correctly solve the problem. This places the spectral method on a stronger theoretical footing as a cross between a heuristic and an exact algorithm. Similar comments apply to envelope reduction. Matrices with dense envelopes are closely related to matrices with the consecutive ones property. Recent work has shown spectral techniques to be better in practice than any existing combinatorial approaches at reducing envelopes [4]. Our result sheds some light on this success. Another way to interpret our result is that we provide an algorithm for C1P that generalizes to become an attractive heuristic in the presence of errors. Designed as decision algorithms for the consecutive ones property, existing combinatorial approaches for C1P break down if there are errors and fail to provide useful approximate orderings. However, our goal here is not to analyze the approach as an approximation algorithm, but rather to prove that it correctly solves error-free problem instances. This paper is organized in the following way. In the next section we introduce the mathematical notation and the results from matrix theory that we will need later. We also describe a spectral heuristic for ordering problems which motivates the remainder of the paper. The theorem that underpins our algorithm is proved in x3, the proof of which requires the use of a classic theorem from matrix analysis. Several additional results in x4 lead us to an algorithm and its analysis in x5. We review the connection to C1P in x6. 2. Mathematical background. 2.1. Notation and Definitions. Matrix concepts are useful because the correlation function defined above can be considered as a real, symmetric matrix. A permutation of the elements corresponds to a symmetric permutation of this matrix, a permutation of the matrix elements formed by permuting the rows and the columns in the same fashion. The question of whether or not the ordering problem is well posed can also be asked as a property of this matrix. Specifically, suppose the matrix has been permuted to reflect a consistent solution to the ordering problem. The off-diagonal matrix entries must now be non-increasing as we move away from the diagonal. More formally, we will say a matrix A is an R-matrix 1 if and only if A is symmetric and a i;j - a i;k for a i;j - a i;k for The diagonal entries of an R-matrix are unspecified. If A can be symmetrically permuted to become an R-matrix, then we say that A is pre-R. Note that pre-R matrices correspond precisely to well-posed ordering problems. Also, the R-matrix property is preserved if we add a constant to all off-diagonal entries, so we can assume without loss of generality that all off-diagonal values are non-negative. When - is a permutation of the natural numbers ng and x is a column vector, i.e. we will denote by x - the permutation of x by -, i.e. Similarly, A - is the symmetric permutation of A by -, i.e. a - We denote by e the vector whose entries are all 1, by e i the vector consisting of zeros except for a 1 in position i, and by I the identity matrix. A symmetric matrix A is reducible if there exists a permutation - such that where B and C are non-empty square matrices. If no such permutation exists then A is irreducible. If B and C are themselves irreducible, then we refer to them as the irreducible blocks of A. We say that - is an eigenvalue of A if corresponding vector x is an eigenvector. An n \Theta n real, symmetric matrix has n eigenvectors that can be constructed to be pairwise orthogonal, and its eigenvalues are all real. We will assume that the eigenvalues are sorted by increasing value, and refer to them as - i , n. The (algebraic) multiplicity of an eigenvalue - is This class of matrices is named after W. S. Robinson who first defined this property in his work on seriation methods in archaeology [29]. defined as the number of times - occurs as a root in the characteristic polynomial value that occurs only once is called simple; the eigenvector of a simple eigenvalue is unique (up to normalization). We write A - 0 and say A is non-negative if all its elements a i;j are non-negative. A real vector x is monotone if We define the Laplacian of a symmetric matrix A to be a diagonal matrix with d a i;j . The minimum eigenvalue with an eigenvector orthogonal to e (the vector of all ones) is called the Fiedler value and a corresponding eigenvector is called a Fiedler vector 2 . Alternatively, the Fiedler value is given by min and a Fiedler vector is any vector x that achieves this minimum while satisfying these constraints. When A - 0 and irreducible, it is not hard to show that the Fiedler value is the smallest non-zero eigenvalue and a Fiedler vector is any corresponding eigenvector. We will be notationally cavalier and refer to the Fiedler value and vector of A when we really mean those of LA . 2.2. PQ-trees. A PQ-tree is a data structure introduced by Booth and Lueker to efficiently encode a set of related permutations [5]. A PQ-tree over a set is a rooted, ordered tree whose leaves are elements of U and whose internal nodes are distinguished as either P-nodes or Q-nodes. A PQ-tree is proper when the following three conditions hold: 1. Every element u appears precisely once as a leaf. 2. Every P-node has at least two children. 3. Every Q-node has at least three children. Two PQ-trees are said to be equivalent if one can be transformed into the other by applying a sequence of the following two equivalence transformations: 1. Arbitrarily permute the children of a P-node. 2. Reverse the children of a Q-node. Conveniently, the equivalence class represented by a PQ-tree corresponds precisely to the set of permutations consistent with an instance of a seriation problem. In x5 we describe an algorithm which uses Laplacian eigenvectors to construct a PQ-tree for an instance of the seriation problem. 2.3. Motivation for Spectral Methods. With the above definitions we can describe a simple heuristic for the seriation problem that will motivate the remainder of the paper. This heuristic is at the heart of the more complex algorithms we will devise, and underlies many previous applications of spectral algorithms [17]. We begin by constructing a simple penalty function g whose value will be small when closely correlated elements are close to each other. We define Unfortunately, minimizing g is NP-hard due to the discrete nature of the permutation [13]. Instead we approximate it by a function h of continuous variables x i that we can minimize and that maintains much of the structure of g. We define . Note that h does not have a unique minimizer, since its value does not change if we add a constant to each x component. To avoid this ambiguity, we need to add a constraint like We still have a trivial solution 2 This is in recognition of the work of Miroslav Fiedler [10, 11]. when all the x i 's are zero, so we need a second constraint like 1. The resulting minimization problem is now well defined. Minimize (1) subject to: The solution to this continuous problem can be used as a heuristic for sequencing. Merely construct the solution vector x, sort the elements x i and sequence based upon their sorted order. One reason this heuristic is attractive is that the minimization problem has an elegant solution. We can rewrite h(x) as x T L F x where is the correlation matrix. The constraints require that x be a unit vector orthogonal to e, and since LA is symmetric, all other eigenvectors satisfy the constraints. Consequently, a solution to the constrained minimization problem is just a Fiedler vector. Even if the problem is not well posed, sorting the entries of the Fiedler vector generates an ordering that tries to keep highly correlated elements near each other. As mentioned above, this technique is being used for a variety of sequencing problems [4, 15, 17]. The algorithm we describe in the remainder of the paper is based upon this idea. However, when we encounter ties in entries of the Fiedler vector, we need to recurse on the subproblem encompassing the tied values. In this way, we are able to find all permutations which make a pre-R matrix into an R matrix. 3. The key theorem. Our main result is that a modification of the simple heuristic presented in x2.3 is actually an algorithm for well-posed instances of the seriation problem. Completely proving this will require us to deal with the special cases of multiple Fiedler vectors and ties within the Fiedler vector. The cornerstone of our analysis is a classical result in matrix theory due to Perron and Frobenius [27]. The particular formulation below can be found on page 46 of [30]. Theorem 3.1 (Perron-Frobenius). Let M be a real, non-negative matrix. If we define 1. ae(M) is an eigenvalue of M , and 2. there is a vector x - 0 such that We are now ready to state and prove our main theorem. Theorem 3.2. If A is an R-matrix then it has a monotone Fiedler vector. Proof. Our proof uses the Perron-Frobenius Theorem 3.1. The non-negative vector in that theorem will consist of differences between neighboring entries in the Fiedler vector of the Laplacian of A. First define the matrix S 2 IR (n\Gamma1)\Thetan as . 0 Note that for any vector x, . 0 It is easy to verify that 1 . We define g. We now show that Sx is an eigenvector of MA if and only if x is an eigenvector of LA and x 6= ffe. The transformation from the second to the third line follows from LA holds between all the above equations, so - is an eigenvalue for both LA and MA for eigenvectors of LA other than e. Hence the eigenvalues of MA are the same as the eigenvalues of LA with the zero eigenvalue removed, and the eigenvectors of MA are differences between neighboring entries of the corresponding eigenvectors of LA . It is easily seen that (SLA ) (a i;k \gamma a i+1;k Since, by assumption, A is an R-matrix, a i;k - a i+1;k for j. For i ? j we can use the fact that (l i;k \gamma l i+1;k Again, from the R-matrix property we conclude that Consequently, all the off-diagonal elements in MA are non-positive. Now let fi be a value greater than are the eigenvalues of MA . Then ~ non-negative with eigenvalues ~ MA and MA share the same set of eigenvectors. By Theorem 3.1, there exists a non-negative eigenvector y of ~ MA corresponding to the largest eigenvalue of ~ MA . But y is also an eigenvector of MA corresponding to MA 's smallest eigenvalue. And this is just Sx, where x is a Fiedler vector of LA . Since non-negative, the corresponding Fiedler vector of LA is non-decreasing and the theorem follows. (Note that since the sign of an eigenvector is unspecified, the Fiedler vector could also be non-increasing.) Theorem 3.3. Let A be a pre-R matrix with a simple Fiedler value and a Fiedler vector with no repeated values. Let - 1 (respectively - 2 ) be the permutation induced by sorting the values in the Fiedler vector in increasing (decreasing) order. Then A - 1 and A - 2 are R-matrices, and no other permutations of A produce R-matrices. Proof. First note that since the Fiedler value is simple, the Fiedler vector is unique up to a multiplicative constant. Next observe that if x is the Fiedler vector of A, then x - is the Fiedler vector of A - . So applying a permutation to A merely changes the order of the entries in the Fiedler vector. Now let - be a permutation such that A - is an R-matrix. By Theorem 3.2 x - is monotone since x is the only Fiedler vector. Since x has no repeated values, - must be either - 1 or - 2 . Theorem 3.3 provides the essence of our algorithm for the seriation problem, but it is too restrictive as the Fiedler value must be simple and contain no repeated values. We will show how to remove these limitations in the next section. 4. Removing the restrictions. Several observations about the seriation problem will simplify our analysis. First note that if we add a constant to all the correlation values the set of solutions is unchanged. Consequently, we can assume without loss of generality that the smallest value of the correlation function is zero. Note that subtracting the smallest value from all correlation values does not change whether or not the matrix is pre-R. In our algebraic formulation this translates into the following. Lemma 4.1. Let A be a symmetric matrix and let - real ff. A vector x is a Fiedler vector of A iff x is a Fiedler vector of - A. So without loss of generality we can assume that the smallest off-diagonal entry of A is zero. Proof. By the definition of a Laplacian it follows that L - where n is the dimension of A. Then L - but for any other eigenvector x of LA , That is, the eigenvalues are simply shifted down by ffn while the eigenvectors are preserved. This will justify the first step of our algorithm, which subtracts the value of the smallest correlation from every correlation. Accordingly, we now make the assumption that our pre-R matrix has smallest off diagonal entry of zero. Next observe that if A is reducible then the seriation problem can be decoupled. The irreducible blocks of the matrix correspond to connected components in the graph of the nonzero values of the correlation function. We can solve the subproblems induced by each of these connected components, and link the pieces together in an arbitrary order. More formally, we have the following lemma. Lemma 4.2. Let A i , be the irreducible blocks of a pre-R matrix A, and let - i be a permutation of block A i such that the submatrix A - i i is an R-matrix. Then any permutation formed by concatenating the - i 's will make A become an R-matrix. In terms of a PQ-tree, the - i permutations are children of a single P-node. Proof. By Lemma 4.1, we can assume all entries in the irreducible blocks are non- negative. Consequently, the correlation between elements within a block will always be at least as strong as the correlation between elements in different blocks. Also, by the definition of irreducibility, each element within a block must have some positive correlation with another element in that block. Hence, any ordering that makes A i an R-matrix must not interleave elements between different irreducible blocks. As long as the blocks themselves are ordered to be R-matrices, any ordering of blocks will make A an R-matrix since correlations across blocks are all identical. With these preliminaries, we will now assume that the smallest off-diagonal value is zero and that the matrix is irreducible. As the following three lemmas and theorem show, this is sufficient to ensure that the Fiedler vector is unique up to a multiplicative constant. Lemma 4.3. Let A be an n \Theta n R-matrix with a monotone Fiedler vector x. If is a maximal interval such that x Proof. We can without loss of generality assume x is non-decreasing since \Gammax is also a Fiedler vector. We will show that a r;k = a s;k for all since A is an R-matrix then all elements between a r;k and a s;k must also be equal. Consider rows r and s in the equation (l s;k \gamma l r;k )x Since LA is a Laplacian, we know that (l sk \gamma l rk )(x r \gamma x k ) (l s;k \gamma l r;k ) (l s;k \gamma l r;k ) where we have used the fact that x is non-decreasing. Because all terms in the sum are non-negative, all terms must be exactly zero. By assumption, x k 6= x r for and consequently l 2 J and the result follows. The following lemma is essentially a converse of this. Its proof requires detailed algebra, but it is not fundamental to what follows. Consequently, the proof is relegated to the end of this section. Lemma 4.4. Let A be an irreducible n \Theta n R-matrix with a [1; n] is an interval such that a r;k = a s;k for all any Fiedler vector x. Lemma 4.5. Let A be an irreducible R-matrix with a Fiedler vector of A. If is an interval such that x for any Fiedler vector y, y Proof. First apply Lemma 4.3 to conclude that for any k = Now use this in conjunction with Lemma 4.4 to obtain the result. Theorem 4.6. If A is an irreducible R-matrix with a then the Fiedler value - 2 is a simple eigenvalue. Proof. We will assume that - 2 is a repeated eigenvalue and produce a contradic- tion. Let x and y be two linearly independent Fiedler vectors with x non-decreasing. sin(')y, with 0 -. Let ' be the smallest value of ' that makes z Such a ' must exist since x and y are linearly independent. By Lemma 4.5 the indices of any repeated values in x are indices of repeated values in y and z('). Coupled with the monotonicity of x, this implies that z(' ) is monotone. By Lemma 4.5 the indices of any repeated values in z(' ) must be repeated in x which gives the desired contradiction. All that remains is to handle the situation where the Fiedler vector has repeated values. As the following theorem shows, repeated values decouple the problem into pieces that can be solved recursively. Theorem 4.7. Let A be a pre-R matrix with a simple Fiedler value and Fiedler vector x. Suppose there is some repeated value fi in x and define I, J and K to be the indices for which 1. 2. x 3. Then - is an R-matrix ordering for A iff - or its reversal can be expressed as (- is an R-matrix ordering for the submatrix A(J ; J ) of A induced by J , and - i and - k are the restrictions of some R-matrix ordering for A to I and K, respectively. Proof. From Theorem 3.2 we know that for any R-matrix ordering A - , x - is monotone, so elements in I must appear before (after) elements from J and elements from K must appear after (before) elements from J . By Lemma 4.3, we have a for all 2 J . Hence the orderings of elements inside J must be indifferent to the ordering outside of J and vice versa. Consequently, the R-matrix ordering of elements in J depends only of A(J ; J ). Algorithmically, this theorem means that we can break ties in the Fiedler vector by recursing on the submatrix A(J ; J ) where J corresponds to the set of repeated values. The distinct values in the Fiedler vector of A constrain R-matrix orderings, but repeated values need to be handled recursively. In the language of PQ-trees, the distinct values are combined via a Q-node, and the components (subtrees) of the Q-node must then be expanded recursively. Proof of Lemma 4.4. First we recall that the Fiedler value is the value obtained by min a i;j (2) and a Fiedler vector is a vector that achieves this minimum. We note that if we replace A by a matrix that is at least as large on an elementwise comparison then x T LAx cannot decrease for any vector x. We consider A(J ; J ), the diagonal block of A indexed by J . By the definition of an R-matrix, all values in A(J ; J ) must be at least as large as a r;s . However, a r;s must be greater than zero. Otherwise, by the R-matrix property a then by the statement of the theorem a which would make the matrix reducible. The remainder of the proof will proceed in two stages. First we will force all the off-diagonal values in A(J ; J ) to be a r;s and show the result for this modified matrix. We will then extend the result to our original matrix. Stage 1: We define the matrix B to be identical to A outside of B(J ; J ), but all off-diagonal values of B within B(J ; J ) are set to ff = a r;s . It follows from the hypotheses that B is an R-matrix. We define note that, by the R-matrix property, We now define ~ ff)I and consider the eigenvalue equation ~ ~ This matrix has the same eigenvectors as LB with eigenvalues shifted by rows of ~ LB in J are identical. Consequently, either all elements of x in J are equal, or ~ - (which is equivalent to - We will show that irreducibility and a which will complete the proof of Stage 1. We assume look for a contradiction. We introduce a new matrix B as follows ff otherwise. Since B is an R-matrix, - B is at least as large as B elementwise, so - 2 ( - We define the vector - y by and - x to be the unit vector in the direction of - y. We note that - and that We have the following chain of inequalities. The last inequality is strict since - b then we can combine an inequality due to Fiedler [10], l ii ; with the observation that min i l i;i - ffi to obtain - 2 - n . This can only be true if equality holds throughout, implying that But this contradicts (3), so - 2 ff and the proof of Stage 1 is complete. Stage 2: We will now show that A and B have the same Fiedler vectors. Since A is elementwise at least as large as B, for any vector z, z T LA z - z T LB z. From Stage 1 we know that any Fiedler vector of B satisfies x . In this vector, for so the contribution to the sum in (2) from B(J ; J ) is zero. But this contribution will also be zero when applied to A(J ; J ). Since A and B are identical outside of A(J ; J ) and B(J ; J ), we now have that a Fiedler vector of B gives an upper bound for the Fiedler value of A; that is, It follows that the Fiedler vectors of B are also Fiedler vectors of A and vice versa. 5. A spectral algorithm for the seriation problem. We can now bring all the preceding results together to produce an algorithm for well-posed instances of the seriation problem. Specifically, given a well-posed correlation function we will generate all consistent orderings. Given a pre-R matrix, our algorithm constructs a PQ-tree for the set of permutations that produce an R-matrix. Our Spectral-Sort algorithm is presented in Fig. 1. It begins by translating all the correlations so that the smallest is 0. It then separates the irreducible blocks (if there are more than one) into the children of a P-node and recurses. If there is only one such block, it sorts the elements into the children of a Q-node based on their values in a Fielder vector. If there are ties in the entries of the Fiedler vector, the algorithm is invoked recursively. Input: A, an n \Theta n pre-R matrix U , a set of indices for the rows/columns of A Output: T , a PQ-tree that encodes the set of all permutations - such that A - is an R-matrix begin (1) ff := min i6=j a i;j (1) A := A \gamma ffee T := the irreducible blocks of A := the corresponding index sets else (3) else if else Fiedler vector for LA number of distinct values in x indices of elements in x with jth value Fig. 1. Algorithm Spectral-Sort. We now prove that the algorithm is correct. Step (1) is justified by Lemma 4.1, and requires time proportional to the number of nonzeros in the matrix. The identification of irreducible blocks in step (2) can be performed with a breadth-first or depth-first search algorithm, also requiring time proportional to the number of nonze- ros. Combining the permutations of the resulting blocks with a P-node is correct by Lemma 4.2. Step (3) handles the boundary conditions of the recursion, while in step (4) the Fiedler vector is computed and sorted. If there are no repeated elements in the Fiedler vector then the Q-node for the permutation is correct by Theorem 3.3. Steps (3) and (4) are the dominant computational steps and we will discuss their run time below. The recursion in step (5) is justified by Theorem 4.7. Note that this algorithm produces a tree whether A is pre-R or not. To determine whether A is pre-R, simply apply one of the generated permutations. If the result is an R-matrix then all permutations in the PQ-tree will solve the seriation problem, otherwise the problem is not well posed. The most expensive steps in algorithm Spectral-Sort are the generation and sorting of the eigenvector. Since the algorithm can invoke itself recursively, these operations can occur on problems of size n, n if the time for an eigen- calculation on a matrix of size n is T (n), the runtime of algorithm Spectral-Sort is A formal analysis of the complexity of the eigenvector calculation can be simplified by noting that for a Pre-R matrix, all that matters is the dominance relationships between matrix entries. So, without loss of generality, we can assume that all entries are integers less than n 2 . With this observation, it is possible to compute the components of the Fiedler vector to a sufficient precision that the components can be correctly sorted in polynomial time. We now sketch one way this can be done, although we don't recommend this procedure in a real-world implementation. Let - denote a specific eigenvalue of L, in our case the Fiedler value. This can be computed in polynomial time as discussed in [25]. Then we can compute the corresponding eigenvector x symbolically by solving where p(z) is the characteristic polynomial of L. Gaussian elimination over a field is in P [21], so if p(z) is irreducible we obtain a solution x where each component x i is given by a polynomial in z with bounded integer coefficients. We note that letting z be any eigenvalue will force x to be a true eigenvector. If p(z) is reducible, we try the above. If we fail to solve the equation, we will instead find a factorization of p(z) and proceed by replacing p(z) with the factor containing - as a root. This yields a polynomial formula for each x i and we can identify equal elements by e.g. the method in [22]. To decide the order of the remaining components, we evaluate the root - to a sufficient precision and then compute the x i 's numerically and sort. Since - is algebraic, the cannot be arbitrarily close [22] and polynomial precision is sufficient. In practice, eigencalculations are a mainstay of the numerical analysis community. To calculate eigenvectors corresponding to the few highest or lowest eigenvalues (like the Fiedler vector), the method of choice is known as the Lanczos algorithm. This is an iterative algorithm in which the dominant cost in each iteration is a matrix-vector multiplication which requires O(m) time. The algorithm generally converges in many fewer than n iterations, often only O( n) [26]. However, a careful analysis reveals a dependence on the difference between the distinct eigenvalues. 6. The consecutive ones problem. Ordering an R-matrix is closely related to the consecutive ones problem. As mentioned in x1, a (0; 1)-matrix C has the consecutive ones property if there exists a permutation matrix \Pi such that for each column in \PiC, all the ones form a consecutive sequence. 3 A matrix that has this property without any rearrangement (i.e. I) is in Petrie form 4 and is called a P- matrix. Analogous to R-matrices, we say a matrix with the consecutive ones property is pre-P. The consecutive ones problem can be restated as: Given a pre-P matrix C, find a permutation matrix \Pi such that \PiC is a P-matrix. 3 Some authors define this property in terms of rows instead of columns. 4 Sir William M. F. Petrie was an archaeologist who studied mathematical methods for seriation in the 1890's. There is a close relationship between P-matrices and R-matrices. The following results are due to D.G. Kendall and are proved in [19] and [33]. Lemma 6.1. If C is a P-matrix, then is an R-matrix. Lemma 6.2. If C is pre-P and is an R-matrix, then C is a P-matrix. Theorem 6.3. Let C be a pre-P matrix, let A = CC T , and let \Pi be a permutation matrix. Then \PiC is a P-matrix if and only if \PiA\Pi T is an R-matrix. This theorem allows us to use algorithm Spectral-Sort to solve the consecutive ones problem. First construct apply our algorithm to A (note that the elements of A are small non-negative integers). Now apply one of the permutations generated by the algorithm to C. If the result is a P-matrix then all the permutations produce C1P orderings. If not, then C has no C1P orderings. The run time for this technique is not competitive with the linear time algorithm for this problem due to Booth and Lueker [5]. However, unlike their approach, our Spectral-Sort algorithm does not break down in the presence of errors and can instead serve as a heuristic. Several other combinatorial problems have been shown to be equivalent to the consecutive ones problem. Among these are recognizing interval graphs [5, 12] and finding dense envelope orderings of matrices [5]. One generalization of P-matrices is to matrices with unimodal columns (a uni-modal sequence is a sequence that is non-decreasing until it reaches its maximum, then non-increasing). These matrices are called unimodal matrices [32]. Kendall [20] showed that the results 6.1 - 6.3 are also valid for unimodal matrices if the regular matrix product is replaced by the matrix circle product defined by Note that P-matrices are just a special case of unimodal matrices, and that the circle product is equivalent to matrix product for (0; 1)-matrices. Kendall's result implies that our spectral algorithm will correctly identify and order unimodal matrices. Acknowledgements . We are indebted to Robert Leland for innumerable discussions about spectral techniques and to Sorin Istrail for his insights into the consecutive ones problem and his constructive feedback on an earlier version of this paper. We are further indebted to David Greenberg for his experimental testing of our approach on simulated genomic data, and to Nabil Kahale for showing us how to simplify the proof of Theorem 3.2. We also appreciate the highly constructive feedback provided by an anonymous referee. --R Interior point methods in semidefinite programming with applications to combinatorial optimization. A spectral technique for coloring random 3-colorable graphs Graph coloring using eigenvalue decomposition. A spectral algorithm for envelope reduction of sparse matrices. Testing for the consecutive ones property A near optimal algorithm for edge seperators (preliminary version). Recent Results in the Theory of Graph Spectra. Spectra of Graphs: Theory and Application. Lower bounds for the partitioning of graphs. Algebraic connectivity of graphs. A property of eigenvectors of nonnegative symmetric matrices and its application to graph theory. Incidence matrices and interval graphs. An analysis of spectral envelope-reduction via quadratic assignment problems Physical mapping by STS hybridization: Algorithmic strategies and the challenge of software evaluation. Geometric Algorithms and Combinatorial Optimiza- tion Optimal linear labelings and eigenvalues of graphs. Approximate graph coloring by semidefinite program- ming Incidence matrices Abundance matrices and seriation in archaeology. The Design and Analysis of Algorithms. of algebraic numbers. The Laplacian spectrum of graphs. Laplace eigenvalues of graphs - a survey Algebraic complexity of computing polynomial zeros. The Lanczos algorithm with selective orthogonalization. Zur Theorie der Matrizen. Partitioning sparse matrices with eigenvectors of graphs. A method for chronologically ordering archaeological deposits. Matrix Iterative Analysis. Approximation of the consecutive ones matrix augmentation problem. Mathematics in the Archaeological and Historical Sciences Techniques of data analysis and seriation theory. --TR --CTR Antonio Robles-Kelly , Edwin R. Hancock, Graph Edit Distance from Spectral Seriation, IEEE Transactions on Pattern Analysis and Machine Intelligence, v.27 n.3, p.365-378, March 2005 Jordi Petit, Experiments on the minimum linear arrangement problem, Journal of Experimental Algorithmics (JEA), 8, Aristides Gionis , Teija Kujala , Heikki Mannila, Fragments of order, Proceedings of the ninth ACM SIGKDD international conference on Knowledge discovery and data mining, August 24-27, 2003, Washington, D.C. Anil Menon, Aspects of the Binary CMAC: Unimodularity and Probabilistic Reconstruction, Neural Processing Letters, v.22 n.3, p.263-276, December 2005 Antti Ukkonen , Mikael Fortelius , Heikki Mannila, Finding partial orders from unordered 0-1 data, Proceeding of the eleventh ACM SIGKDD international conference on Knowledge discovery in data mining, August 21-24, 2005, Chicago, Illinois, USA Huaijun Qiu , Edwin R. Hancock, Graph matching and clustering using spectral partitions, Pattern Recognition, v.39 n.1, p.22-34, January, 2006 Sebastian Lange , Martin Middendorf, Design Aspects of Multi-level Reconfigurable Architectures, Journal of Signal Processing Systems, v.51 n.1, p.23-37, April 2008

consecutive ones property;eigenvector;analysis of algorithms;fiedler vector;seriation

298560

New Collapse Consequences of NP Having Small Circuits.

We show that if a self-reducible set has polynomial-size circuits, then it is low for the probabilistic class ZPP (NP). As a consequence we get a deeper collapse of the polynomial-time hierarchy PH to ZPP(NP) under the assumption that NP has polynomial-size circuits. This improves on the well-known result in Karp and Lipton [ Proceedings of the 12th ACM Symposium on Theory of Computing, ACM Press, New York, 1980, pp. 302--309] stating a collapse of PH to its second level $\Sigmap_2$ under the same assumption. Furthermore, we derive new collapse consequences under the assumption that complexity classes like UP, FewP, and C=P have polynomial-size circuits. Finally, we investigate the circuit-size complexity of several language classes. In particular, we show that for every fixed polynomial s, there is a set in ZPP(NP) which does not have O(s(n))-size circuits.

Introduction . The question of whether intractable sets can be efficiently decided by non-uniform models of computation has motivated much work in structural complexity theory. In research from the early 1980's to the present, a variety of results has been obtained showing that this is impossible under plausible assumptions (see, e.g., the survey [18]). A typical model for non-uniform computations are circuit families. In the notation of Karp and Lipton [22], sets decidable by polynomial-size circuits are precisely the sets in P/poly, i.e., they are decidable in polynomial time with the help of a polynomial length bounded advice function [32]. Karp and Lipton (together with Sipser) [22] proved that no NP-complete set has polynomial size circuits (in symbols NP 6' P/poly) unless the polynomial time hierarchy collapses to its second level. The proof given in [22] exploits a certain kind of self-reducibility of the well-known NP-complete problem SAT. More generally, it is shown in [8, 7] that every (Turing) self-reducible set in P/poly is low for the second level \Sigma P 2 of the polynomial time hierarchy. Intuitively speaking, a set is low for a relativizable complexity class if it gives no additional power when used as an oracle for that class. In this paper, we show that every self-reducible set in P/poly is even low for the probabilistic class ZPP(NP), meaning that for every oracle A, \Sigma lowness for ZPP(NP) implies lowness for \Sigma P. As a consequence of our lowness result we get a deeper collapse of the polynomial-time hierarchy to ZPP(NP) under the assumption that NP has polynomial-size circuits. At Abteilung f?r Theoretische Informatik, Universit?t Ulm, Oberer Eselsberg, D-89069 Ulm, Germany ([email protected]). y Department of Computer Science, Tokyo Institute of Technology, Meguro-ku, Tokyo 152, Japan ([email protected]). Part of this work has been done while visiting the University of Ulm (supported in part by the guest scientific program of the University of Ulm). J. K - OBLER AND O. WATANABE least in some relativized world, the new collapse level is quite close to optimal: there is an oracle relative to which NP is contained in P/poly but PH does not collapse to P(NP) [17, 39]. We also derive new collapse consequences from the assumption that complexity classes like UP, FewP, and C=P have polynomial-size circuits. Furthermore, our lowness result implies new relativizable collapses for the case that Modm P, PSPACE, or EXP have polynomial-size circuits. As a final application, we derive new circuit- size lower bounds. In particular, it is shown (by relativizing proof techniques) that for every fixed polynomial s, there is a set in ZPP(NP) which does not have O(s(n))-size circuits. This improves on the result of Kannan [21] that for every polynomial s, the class \Sigma P 2 contains such a set. It further follows that in every relativized world, there exist sets in the class ZPEXP(NP) that do not have polynomial-size circuits. It should be noted that there is a non-relativizing proof for a stronger result. As a corollary to the result in [4], which is proved by a non-relativizing technique, it is provable that MA exp " co-MA exp (a subclass of ZPEXP(NP)) contains non P/poly sets [12, 36]. Some explanation of how our work builds on prior techniques is in order. The proof of our lowness result heavily uses the universal hashing technique [13, 34] and builds on ideas from [2, 14, 24]. For the design of a zero error probabilistic algorithm which, with the help of an NP oracle, simulates a given ZPP(NP(A)) computation (where A is a self-reducible set in P/poly) we further make use of the newly defined concept of half-collisions. More precisely, we show how to compute on input 0 n in expected polynomial time a hash family H that can be used to decide all instances of A of length up to n by a strong NP computation. The way H is used to decide (non)membership to A is by checking whether H leads to a half-collision on certain sets. Very recently, Bshouty, Cleve, Gavald'a, Kannan, and Tamon [11] building on a result from [19] have shown that the class of all circuits is exactly learnable in (randomized) expected polynomial time with equivalence queries and the aid of an NP oracle. This immediately implies that for every set A in P/poly an advice function can be computed in FZPP(NP(A)), i.e., by a probabilistic oracle transducer T in expected polynomial time under an oracle in NP(A). More precisely, since the circuit produced by the probabilistic learning algorithm of [11] depends on the outcome of the coin flips, T computes a multi-valued advice function, i.e., on input 0 n , T accepts with probability at least 1/2, and on every accepting path, T outputs some circuit that correctly decides all instances of length n w.r.t. A. Using the technique in [11] we are able to show that every self-reducible set A in P/poly even has an advice function in FZPP(NP). Although this provides a different way to deduce the ZPP(NP) lowness of all self-reducible sets in P/poly, we prefer to give a self-contained proof using the "half-collision technique" that does not rely on the mentioned results in [11, 19]. The paper is organized as follows: Section 2 introduces notation and defines the self-reducibility that we use. In Section 3 we prove the ZPP(NP) lowness of all self- reducible sets in P/poly. In Section 4 we state the collapse consequences, and the new circuit-size lower bounds are derived in Section 5. 2. Preliminaries and notation. All languages are over the binary alphabet 1g. As usual, we denote the lexicographic order on \Sigma by -. The length of a string x 2 \Sigma is denoted by jxj. \Sigma -n (\Sigma !n ) is the set of all strings of length at most n (resp., of length smaller than n). For a language A, A . The cardinality of a finite set A is denoted by jAj. The characteristic function of A is defined as otherwise. For a class C of sets, co-C denotes the class f\Sigma \Gamma A j A 2 Cg. To encode pairs (or tuples) of strings we use a standard polynomial-time computable pairing function denoted by h\Delta; \Deltai whose inverses are also computable in polynomial time. Where intent is clear we in place of f(hx denotes the set of non-negative integers. Throughout the paper, the base of log is 2. The textbooks [9, 10, 25, 31, 33] can be consulted for the standard notations used in the paper and for basic results in complexity theory. For definitions of probabilistic complexity classes like ZPP see also [15]. An machine M is a polynomial-time nondeterministic Turing machine. We assume that each computation path of M on a given input x either accepts, rejects, or outputs "?". M accepts on input x, if M performs at least one accepting computation, otherwise M rejects x. M strongly accepts (strongly rejects) x [26] if ffl there is at least one accepting (resp., rejecting) computation path and ffl there are no rejecting (resp., accepting) computation paths. If M strongly accepts or strongly rejects x, M is said to perform a strong computation on input x. An NP machine that on every input performs a strong computation is called a strong NP machine. It is well known that exactly the sets in NP " co-NP are accepted by strong NP machines [26]. Next we define the kind of self-reducibility that we use in this paper. Definition 2.1. Let - be an irreflexive and transitive order relation on \Sigma . A sequence of strings is called a -chain (of length k) from x 0 to x k if Relation - is called length checkable if there is a polynomial q such that 1. for all x; y 2 \Sigma , x - y implies jyj - q(jxj), 2. the language fhx; there is a -chain of length k from x to yg is in NP. Definition 2.2. A set A is self-reducible, if there is a polynomial-time oracle machine M self and a length checkable order relation - such that on any input x, M self queries the oracle only about strings y OE x. It is straightforward to check that the polynomially related self-reducible sets introduced by Ko [23] as well as the length-decreasing and word-decreasing self-reducible sets of Balc'azar [6] are self-reducible in our sense. Furthermore, it is well-known (see, for example, [9, 6, 29]) that complexity classes like NP, \Sigma P PSPACE, and EXP have many-one complete self-reducible sets. Karp and Lipton [22] introduced the notion of advice functions in order to characterize non-uniform complexity classes. A function h : N ! \Sigma is called a polynomial- length function if for some polynomial p and for all n - 0, p(n). For a class C of sets, let C=poly be the class of sets A such that there is a set I 2 C and a 4 J. K - OBLER AND O. WATANABE polynomial-length function h such that for all n and for all x 2 \Sigma -n , Function h is called an advice function for A, whereas I is the corresponding interpreter set. In this paper we will heavily make use of the "hashing technique" which has been very fruitful in complexity theory. Here we review some notations and facts about hash families. We also extend the notion of "collision" by introducing the concept of a "half-collision" which is central to our proof technique. Sipser [34] used universal hashing, originally invented by Carter and Wegman [13], to decide (probabilistically) whether a finite set X is large or small. A linear function h from \Sigma m to \Sigma k is given by a Boolean (k; m)-matrix (a ij ) and maps any string x to a string is the inner product of the i-th row a i and x. h be a linear hash function from \Sigma m to \Sigma k . Then we say that x has a collision on Y w.r.t. h if there exists a string y 2 Y , different from x, such that generally, if X is a subset of \Sigma m and H is a family l ) of linear hash functions from \Sigma m to \Sigma k , then we say that X has a collision on Y w.r.t. H (Collision(X; Y; H) for short) if there is some x 2 X that has a collision on Y w.r.t. every h i in H. That is, and for all If X has a collision on itself w.r.t. H, we simply say that X has a collision w.r.t. H. Next we extend the notion of "collision" in the following way. For any X and Y family l ) of linear hash functions, we say that X has a half-collision on Y w.r.t. H (Half-Collision(X; Y; H) for short) if there is some x 2 X that has a collision on Y w.r.t. at least dl=2e many of the hash functions h i in H. That is, Half-Collision(X; and An important relationship between collisions and half-collisions is the following one: If X has a collision w.r.t. H on must have a half-collision w.r.t. H either on Y 1 or on Y 2 . Note that the predicate Collision(X; Y; H) can be decided in NP provided that membership in X and Y can be tested in NP. More precisely, the language fhv; Hi j (as well as the set fhv; Hi j Half-Collision(X to NP, if the sets X v and Y v are succinctly represented in such a way that the languages are in NP. We denote the set of all families l ) of l linear hash functions from \Sigma m to \Sigma k by H(l; m; k). The following theorem is proved by a pigeon-hole argument. It says that every sufficiently large set must have a collision w.r.t. any hash family. Theorem 2.3. [34] For any hash family H 2 H(l; m; k) and any set X ' \Sigma m of cardinality jXj ? l must have a collision w.r.t. H. On the other hand, we get from the next theorem (called Coding Lemma in [34]) an upper bound on the collision probability for sufficiently small sets. Theorem 2.4. [34] Let X ' \Sigma m be a set of cardinality at most 2 k\Gamma1 . If we choose a hash family H uniformly at random from H(k; m; k), then the probability that X has a collision w.r.t. H is at most 1=2. We will also make use of the following extension of Theorem 2.4 which can be proved along the same lines. Theorem 2.5. Let X ' \Sigma m be a set of cardinality at most 2 k\Gammas . If we choose a family H uniformly at random from H(l; m; k), then the probability that X has a collision w.r.t. H is at most 2 k\Gammas(l+1) . Gavald'a [14] extended Sipser's Coding Lemma (Theorem 2.4) to the case of a collection C of exponentially many sets. The following theorem has a similar flavor. Theorem 2.6. Let C be a collection of at most 2 n subsets of \Sigma m , each of which has cardinality at most 2 k\Gammas . If we choose a hash family H uniformly at random from then the probability that some X 2 C has a collision w.r.t. H is at most Proof. By Theorem 2.5, we have that for every fixed X 2 C, the probability that it has a collision w.r.t. a randomly chosen hash family H 2 H(l; m; k) is at most 2 k\Gammas(l+1) . Hence, the probability that there exists such a set X 2 C is at most In this paper we make use of a corresponding result for the case of half-collisions. Theorem 2.7. Let X ' \Sigma m and let C be a collection of at most 2 n subsets of , each of which has cardinality at most 2 k\Gammas\Gamma2 . If we choose a hash family H uniformly at random from H(l; m; k), then the probability that X has a half-collision on some Y 2 C w.r.t. H is at most jXj \Delta 2 n\Gammasl=2 . Proof. For every fixed Y 2 C and every fixed x 2 X, the probability that x has a collision on Y w.r.t. a randomly chosen h is at most 2 \Gammas\Gamma2 . Hence, the probability that x has a collision on Y w.r.t. at least half of the functions in a randomly chosen hash family H 2 H(l; m; k) is at most l l That is, the probability that x has a half-collision on Y w.r.t. a randomly chosen hash family H is bounded by 2 \Gammasl=2 . Hence, the probability that there exists a Y 2 C and an x 2 X such that x has a half-collision on Y w.r.t. H is at most jXj \Delta 2 n\Gammasl=2 . 3. Lowness of self-reducible sets in P/poly. In this section, we show that every self-reducible set A in (NP " co-NP)=poly is low for ZPP(NP). Let I 2 NP " co-NP 6 J. K - OBLER AND O. WATANABE be an interpreter set and h be an advice function for A. We construct a probabilistic algorithm T and an NP oracle O having the following two properties: a) The expected running time of T is polynomially bounded. b) On every computation path on input 0 n , T with oracle O outputs some information that can be used to determine the membership to A of any x up to length n by some strong NP computation (in the sense of [26]). Using these properties, we can prove the lowness of A for ZPP(NP) as follows: In order to simulate any NP(A) computation, we first precompute the above mentioned information for A (up to some length) by T O , and then by using this information, we can simulate the NP(A) computation by some NP(NP " co-NP) computation. Note that the precomputation (performed by T O ) can be done in ZPP(NP), and since the remaining computation can be done in NP. Hence, which implies further that ZPP(NP(A)) ' ZPP(ZPP(NP)) (= ZPP(NP) [41]). We will now make the term "information" precise. For this, we need some additional notation. Let the self-reducibility of A be witnessed by a polynomial-time oracle machine M self , a length checkable order relation -, and a polynomial q. We assume that fixed polynomial p ? 0. In the following, we fix n and consider instances of length up to q(n) as well as advice strings of length exactly p(n). ffl A sample is a sequence hx of pairs, where the x i 's are instances of length up to q(n) and b ffl For any sample Consistent(S) be the set of all advice strings w that are consistent with S, i.e. The cardinality of Consistent(S) is denoted by c(S). ffl For any sample S and any instance x, let Accept(x; S) (resp., Reject(x; S)) be the set of all consistent advice strings that accept x (resp., reject x): and ffl Let Correct(x; S) be the set fw 2 Consistent(S) j I(x; of consistent advice strings that decide x correctly, and let Incorrect(x; S) be the complementary set fw 2 Consistent(S) j I(x; w) 6= A(x)g. Note that the sets Accept(x; S) and Reject(x; S) (as well as Correct(x; S) and Incorrect(x; S)) form a partition of the set Consistent(S), and that The above condition b) can now be precisely stated as follows: b) On every computation path on input 0 n , T O outputs a pair hS; Hi consisting of a sample S and a linear hash family H such that for all x up to length n, Consistent(S) has a half-collision w.r.t. H on Correct(x; S), but not on Incorrect(x; S). Once we have a pair hS; Hi satisfying condition b), we can determine whether an instance x of length up to n is in A by simply checking whether Consistent(x; S) has a half-collision w.r.t. H on Accept(x; S) or on Reject(x; S). Since condition b) guarantees that the half-collision can always be found, this checking can be done by a strong NP computation. Let us now prove our main lemma. Lemma 3.1. For any self-reducible set A in (NP " co-NP)=poly, there exist a probabilistic transducer T and an oracle O in NP satisfying the above two conditions. Proof. We use the notation introduced so far. Recall that q(n) is a length bound on the queries occuring in the self-reduction tree produced by M self on any instance of length n and that p(n) is the advice length for the set of all instances of length up to q(n). Let l be the polynomial defined as 1). Further, we denote by \Sigma -n the set fy j 9x 2 \Sigma -n ; y - xg. Then it is clear that \Sigma -n ' \Sigma -n ' \Sigma -q(n) . A description of T is given below. input loop randomly from H(l(n); p(n); k), has a collision w.r.t. H k g if there exists an x 2 \Sigma -n such that Consistent(S) has a half-collision on Incorrect(x; S) w.r.t. H kmax then use oracle O to find such a string x and to determine A(x) else exit(loop) end loop output Starting with the empty sample, T enters the main loop. During each execution of the loop, T first randomly guesses a series of p(n) many hash families computes the integer kmax as the maximum p(n)g such that Consistent(S) has a collision w.r.t. H k . Notice that by a padding trick we can assume that c(S) is always larger than 2l(n), implying that Consistent(S) must have a collision w.r.t. H 1 . Since, in particular, Consistent(S) has a collision w.r.t. H kmax , it follows that for every instance x 2 \Sigma -n , Consistent(S) has a half-collision w.r.t. H kmax on either Correct(x; S) or Incorrect(x; S). If there exists a string x 2 \Sigma -n such that Consistent(S) has a half-collision on Incorrect(x; S) w.r.t. H kmax , then this string is added to the sample S and T continues executing the loop. (We will describe below how T uses the NP oracle O to find x in this case.) Otherwise, 8 J. K - OBLER AND O. WATANABE the pair hS; H kmax i fulfills the properties stated in condition b) and T halts. We now show that the expected running time of T is polynomially bounded. Since the initial size of Consistent(S) is 2 p(n) , and since Consistent(S) never becomes empty, it suffices to prove that for some polynomial r, T eliminates in each single execution of the main loop with probability at least 1=r(n) at least an 1=r(n)-fraction of the circuits in Consistent(S). In fact, we will show that each single extension of S by a reduces the size of Consistent(S) with probability at least by a factor smaller than can only perform more than 2 7 l(n)p(n) loop iterations, if during some iteration of the main loop T extends S by a pair hx; A(x)i which does not shrink the size of Consistent(S) by a factor smaller than the probability for this event is bounded by 2 7 l(n)p(n) Let S be a sample and let kmax be the corresponding integer as determined by T during some specific execution of the loop. We first derive a lower bound for kmax . be the smallest integer k - 1 such that c(S) - l(n)2 k+1 . Since either p(n) or Consistent(S) does not have a collision w.r.t. the hash family H kmax 1), we have (using Theorem 2.3) that c(S) - l(n)2 kmax +1 . Hence, Since T expands S only by strings x 2 \Sigma -n such that Consistent(S) has a half-collision on Incorrect(x; S) w.r.t. H kmax , and since Consistent(S#hx; the probability that the size of Consistent(S) does not decrease by a factor smaller than bounded by the probability that, w.r.t. H kmax , Consistent(S) has a half-collision on some set Incorrect(x; S) of size at most c(S)=2 7 l(n). Let it follows from Theorem 2.7 that the probability of Consistent(S) having a half-collision on some w.r.t. a uniformly at random chosen hash family H 2 H(l(n); at most Thus the probability that for some k - 0, Consistent(S) has a half-collision w.r.t. H k0+k on some set Incorrect(x; S) which is of size at most c(S)=2 7 l(n) is bounded by We finally show how T determines an instance x 2 \Sigma -n (if it exists) such that Consistent(S) has a half-collision on Incorrect(x; S) w.r.t. H kmax . Intuitively, we use the self-reducibility of A to test the "correctness" w.r.t. A of the "program" hS; H kmax i, where we say that ffl a pair hS; Hi accepts an instance x if Consistent(S) has a half-collision on ffl hS; Hi rejects x if Consistent(S) has a half-collision on Reject(x; S) w.r.t. H. Notice that an (incorrect) program might accept and at the same time reject an instance. The main idea to find out whether hS; H kmax i is incorrect on some instance (meaning that w.r.t. H kmax Consistent(S) has a half-collision on Incorrect(x; S)) is to test whether the program hS; H kmax i is in accordance with the output of M self when the oracle queries of M self are answered according to the program To be more precise, consider the NP set Hi j there is a computation path - of M self on input z fulfilling the following properties: - if a query q is answered 'yes', then hS; Hi accepts q, - if a query q is answered 'no', then hS; Hi rejects q, if - is accepting, then hS; Hi rejects z, and if - is rejecting, then hS; Hi accepts z g. Then, as shown by the next claim, the correctness of hS; H kmax i on an instance z can be decided by asking whether hz; belongs to B, provided that hS; H kmax i is correct on all potential queries of M self on input z. Claim. Assume that hS; H kmax i is correct on all y OE z. Then hS; H kmax i is incorrect on z if and only if hz; belongs to B. Proof. Using the fact that for every instance x 2 \Sigma -n , Consistent(S) has a half- collision w.r.t. H kmax on either Correct(x; S) or Incorrect(x; S), it is easy to see that if is incorrect on z, then the computation path - followed by M self (z) under oracle A witnesses hz; B. For the converse, assume that hz; belongs to B and let - be a computation path witnessing this fact. Note that all queries q on - are answered correctly w.r.t. A, since otherwise hS; H kmax i were incorrect on q OE z. Hence, - is the path followed by M self (z) under oracle A and therefore decides z correctly. On the other hand, since - witnesses hz; indeed is incorrect on z. Now we can define the oracle set O as C \Phi D, where Hi j there is a -chain of length (at least) k from some string y 2 \Sigma -n to some string z - x such that hz; and Hi j there is an accepting computation path - of M self on input x such that any query q is only answered 'yes' (`no') if Consistent(S) has a half-collision on Accept(q; S) (resp., Reject(q; S)) w.r.t. H g. Note that the proof of the claim above also shows that for any z 2 \Sigma -n such that is correct on all y OE z, z 2 A if and only if hz; to D. Now we can complete the description of T . T first asks whether the string belongs to C. It is clear that a negative answer implies that is correct on \Sigma -n . Otherwise, by asking queries of the form computes by binary search i max as the maximum value belongs to C (a similar idea is used OBLER AND O. WATANABE input loop randomly from H(l(n); p(n); k), has a collision w.r.t. H k g if else exit(loop) end loop output in [27]). Knowing i determines the lexicographically smallest string xmin such that h0 is in C. Since hq; holds for all instances q OE xmin , it follows inductively from the claim that hS; H kmax i is correct on all q OE xmin . Hence, must be incorrect on xmin , and furthermore, T can determine the membership of xmin to A by asking whether the string hx min ; belongs to D. Theorem 3.2. Every self-reducible set A in the class (NP " co-NP)=poly is low for ZPP(NP). Proof. We first show that NP(A) ' ZPP(NP). Let L be a set in NP(A), and let M be a deterministic polynomial-time oracle machine such that for some polynomial t, Let s(n) be a polynomial bounding the length of all oracle queries of M on some input hx; yi where x is of length n. Then L can be accepted by a probabilistic oracle machine N using the following NP oracle O Hi j there is a y 2 \Sigma t(jxj) such that M on input hx; yi has an accepting path - on which each query q is answered 'yes' (`no') only if Consistent(S) has a half-collision on Accept(q; S) (resp., Reject(q; S)) w.r.t. H g. Here is how N accepts L. On input x, N first simulates T on input 0 s(jxj) to compute a pair hS; H kmax i as described above (T asks questions to some NP oracle O). Then N asks the query hx; O 0 to find out whether x is in L. This proves that NP(A) ' ZPP(NP). Since via a proof that relativizes, it follows that ZPP(NP(A)) is also contained in ZPP(NP), showing that A is low for ZPP(NP). 4. Collapse consequences. As a direct consequence of Theorem 3.2 we get an improvement of Karp, Lipton, and Sipser's result [22] that NP is not contained in P/poly unless the polynomial-time hierarchy collapses to \Sigma P . Corollary 4.1. If NP is contained in (NP " co-NP)=poly then the polynomial-time hierarchy collapses to ZPP(NP). Proof. Since the NP-complete set SAT is self-reducible, the assumption that NP is contained in (NP " co-NP)=poly implies that SAT is low for ZPP(NP), and hence the polynomial-time hierarchy collapses to ZPP(NP). The collapse of the polynomial-time hierarchy deduced in Corollary 4.1 is quite close to optimal, at least in some relativized world [17, 39]: there is an oracle relative to which NP is contained in P/poly but the polynomial-time hierarchy does not collapse to P(NP). In the rest of this section we report some other interesting collapses which can be easily derived using (by now) standard techniques, and which have also been pointed out independently by several researchers to the second author. First, it is straightforward to check that Theorem 3.2 relativizes: For any oracle B, if A is a self-reducible set in the class (NP(B) " co-NP(B))=poly, then NP(A) is contained in ZPP(NP(B)). Consequently, Theorem 3.2 generalizes to the following result. Theorem 4.2. If A is a self-reducible set in the class (\Sigma P )=poly, then As a direct consequence of Theorem 4.2 we get an improvement of results in [1, 20] stating (for k is not contained in (\Sigma P )=poly unless the polynomial-time hierarchy collapses to \Sigma P . Corollary 4.3. Let k - 1. If \Sigma P k is contained in (\Sigma P )=poly, then the polynomial-time hierarchy collapses to ZPP(\Sigma P Proof. Since \Sigma P contains complete self-reducible languages, the assumption that k is contained in (\Sigma P )=poly implies that \Sigma P Yap [40] proved that \Pi P k is not contained in \Sigma P =poly unless the polynomial-time hierarchy collapses to \Sigma P k+2 . As a further consequence of Theorem 4.2 we get the following improvement of Yap's result. Corollary 4.4. For k - 1, if \Pi P =poly, then Proof. The assumption that \Pi P k is contained in \Sigma P =poly implies that \Sigma P k+1 is contained in \Sigma P =poly ' (\Sigma P )=poly. Hence we can apply Corollary 4.3. As corollaries to Theorem 4.2, we also have similar collapse results for many other complexity classes. What follows are some typical examples. Corollary 4.5. For K 2 co-NP)=poly then K is low for ZPP(NP). Proof. It is well-known that for every set A in UP (FewP), the left set of A [30] is word-decreasing self-reducible and in UP (resp., FewP). Thus, under the assumption that UP ' (NP " co-NP)=poly (resp., FewP ' (NP " co-NP)=poly) it follows by Theorem 3.2 that the left set of A (and since A is polynomial-time many-one reducible to its left set, also is low for ZPP(NP). Corollary 4.6. For every k - 1, if C=P ' (\Sigma P )=poly then Proof. First, since C=P has complete word-decreasing self-reducible languages )=poly implies C=P ' ZPP(\Sigma P OBLER AND O. WATANABE )=poly implies PH ' (\Sigma P k )=poly and therefore PH collapses to ZPP(\Sigma P k ) by Corollary 4.3. Finally, since C=P(PH) ' BPP(C= P) [37], it follows that C=P(PH) ' PH, and since [38]), we get inductively that CH ' PH (' ZPP(\Sigma P Corollary 4.7. Let K 2. If for some k - 1, )=poly, then K ' PH and PH collapses to ZPP(\Sigma P Proof. The proof for K 2 fEXP;PSPACEg is immediate from Theorem 4.2 since PSPACE has complete (length-decreasing) self-reducible languages, and since EXP has complete (word-decreasing) self-reducible languages [6]. The proof for K 2 is analogous to the one of Corollary 4.6 using the fact that ModmP has complete word-decreasing self-reducible languages [29], and that PH ' BPP(Modm P) [37, 35]. Since our proof technique is relativizable, the above results hold for every relativized world. On the other hand, it is known that for some classes stronger collapse consequences can be obtained by using non-relativizable arguments. Theorem 4.8. [28, 4, 3] For K 2 fPP; ModmP;PSPACE;EXPg, if K ' P/poly then K ' MA. Harry Buhrman pointed out to us that Corollary 4.7 can also be derived from Theorem 4.8. 5. Circuit complexity. Kannan [21] proved that for every fixed polynomial s, there is a set in \Sigma P which cannot be decided by circuits of size s(n). Using a padding argument, he obtained the existence of sets in NEXP(NP) " co-NEXP(NP) not having polynomial-size circuits. Theorem 5.1. [21] 1. For every polynomial s, there is a set in \Sigma P 2 that does not have circuits of size s(n). 2. For every increasing time-constructible super-polynomial function f(n), there is a set in NTIME[f(n)](NP)"co-NTIME[f(n)](NP) that does not have polynomial size circuits. As an application of our results in Section 3, we can improve Kannan's results in every relativized world from the class \Sigma P 2 to ZPP(NP), and from the class " co-NTIME[f(n)](NP) to ZPTIME[f(n)](NP), respectively. Here ZPTIME[f(n)](NP) denotes the class of all sets that are accepted by some zero error probabilistic machine in expected running time O(f(n)) relative to some NP oracle. Note that for all sets in the class P/poly we may fix the interpreter set to some appropriate one in P. Let I univ denote such a fixed interpreter set. Furthermore, P/poly remains the same class, if we relax the notion of an advice function h (w.r.t. I univ ) as follows: For every x, I univ (x; h(jxj)), i.e., h(n) has to decide correctly only A =n (instead of A -n ). A sequence of circuits Cn , n - 0, is called a circuit family for A, if for every n - 0, Cn has n input gates, and for all n-bit strings x 1 It is well-known (see, e.g., [9]) that I univ can be chosen in such a way that advice length and circuit size (i.e., number of gates) are polynomially related to each other. More precisely, we can assume that there is a polynomial p such that the following holds for every set A. ffl If h is an advice function for A w.r.t. I univ , then there exists a circuit family Cn , n - 0, for A of size jCn j - p(n ffl If Cn , n - 0, is a circuit family for A, then there exists an advice function h for A w.r.t. I univ of length jh(n)j - p(jC n j). Moreover, we can assume that for every polynomial-time interpreter set I there is a constant c I such that if h is an advice function for A w.r.t. I, then there exists an advice function h 0 for A w.r.t. I univ of length jh 0 (n)j I for all n. The following lemma is obtained by a direct diagonalization (cf. the corresponding result in [21]). A set S is called -printable (see [16]) if there is a polynomial-time oracle transducer T and an oracle set A 2 C such that on any input 0 n , T A outputs a list of all strings in S -n . Lemma 5.2. For every fixed polynomial s, there is a \Delta P 3 -printable set A such that every advice function h for A is of length jh(n)j - s(n), for almost all n. Proof. For a given n, be the sequence of strings of length n, enumerated in lexicographic order. Consider the two sets Have-Advice and Find -A defined as follows: Have-Advice , 9 a j+1 \Delta \Delta \Delta a Since there are only 2 strings w in \Sigma !s(n) , at least one pair of the form hn; a 1 \Delta \Delta \Delta a s(n) i is not contained in Have-Advice (provided that s(n) - 2 n ). Let ff n denote the lexicographically smallest such pair hn; a 1 \Delta \Delta \Delta a s(n) i, i.e., there is no advice of length smaller than s(n) that accepts the strings x according to A as the set of all strings x i (jx n) such that 1 - i - s(n) - 2 n and the ith bit of ff n (i.e., a i ) is 1. By a binary search using oracle Find -A, ff n is computable in polynomial time. Since Have-Advice is in NP and thus Find -A is in NP(NP), it follows that A is P(NP(NP))-printable. Since furthermore, for almost all n, A =n has no advice of length smaller than s(n), the lemma follows. Corollary 5.3. For every fixed polynomial s, there is a set A in ZPP(NP) that does not have circuits of size s(n). Proof. If NP does not have polynomial-size circuits, then we can take Otherwise, by Corollary 4.1, and thus the theorem easily follows from Lemma 5.2. Corollary 5.4. Let f be an increasing, time-constructible, super-polynomial function. Then ZPTIME[f(n)](NP) contains a set A that does not have polynomial-size circuits. 14 J. K - OBLER AND O. WATANABE Proof. If NP does not have polynomial-size circuits, then we can take Otherwise, by Corollary 4.1, and thus it follows from Lemma 5.2 that there is a set B in ZPTIME[n k ](NP) such that every advice function h for B is of length jh(n)j - n for almost all n. By the proof technique of Lemma 5.2, we can assume that in all length n strings of B, 1's only occur at the O(log n) rightmost positions. Now consider the following set (where n denotes jxj) and the interpreter set Clearly, A belongs to ZPTIME[f(n)](NP) and I belongs to P. Furthermore, if h is an advice function for A, then we have for every y of the form 0 bf(n) 1=k c\Gamman x, that where h 0 (n) is a suitable advice function of length jh 0 (n)j I . Thus, it follows for almost all n that This shows that the length of h is super-polynomial. Corollary 5.5. In every relativized world, ZPEXP(NP) contains sets that do not have polynomial-size circuits. We remark that the above results are proved by relativizable arguments. On the other hand, Harry Buhrman [12] and independently Thomas Thierauf [36] pointed out to us that Theorem 4.8 (which is proved by a non-relativizable proof technique) can be used to show that MA exp " co-MA exp contains non P/poly sets. Here, MA exp denotes the exponential-time version of Babai's class MA [5]. That is, MA where a language L is in MA[f(n)], if there exists a set B 2 DTIME[O(n)] such that for all x of length n, where z is chosen uniformly at random from \Sigma f(n) . Corollary 5.6. [12, 36] MA exp " co-MA exp contains sets that do not have polynomial size circuits. Since there exist recursive oracles relative to which all sets in EXP(NP) have polynomial size circuits [39, 17], it is not possible to extend Corollary 5.5 by relativizing techniques to the class EXP(NP). 6. Concluding remarks. An interesting question concerning complexity classes C that are known to be not contained in P/poly but are not known to have complete sets is whether the existence of sets in C \Gamma P/poly can be constructively shown. For example, by Corollary 5.5 we know that the class ZPEXP(NP) contains sets that do not have polynomial-size circuits. But we were not able to give a constructive proof of this fact. To the best of our knowledge, no explicit set is known even in Acknowledgments For helpful discussions and suggestions regarding this work we are very grateful to H. Buhrman, R. Gavald'a, L. Hemaspaandra, M. Ogihara, U. Sch-oning, R. Schuler, and T. Thierauf. We like to thank H. Buhrman, L. Hemaspaandra, and M. Ogihara for permitting us to include their observations in the paper. --R On hiding information from an oracle Queries and concept learning Arithmetization: A new method in structural complexity a randomized proof system and a hierarchy of complexity classes Structural Complexity Theory I Introduction to the Theory of Complexity Oracles and queries that are sufficient for exact learning Universal classes of hash functions Bounding the complexity of advice functions Computational complexity of probabilistic complexity classes Computation times of NP sets of different densities On relativized exponential and probabilistic complexity classes How hard are sparse sets? Random generation of combinatorial structures from a uniform distribution Some connections between nonuniform and uniform complexity classes Journal of Computer and System Sciences Strong nondeterministic polynomial-time reducibilities Algebraic methods for interactive proof systems On sparse hard sets for counting classes On polynomial-time bounded truth-table reducibility of NP sets to sparse sets Computational Complexity On simultaneous resource bounds A complexity theoretic approach to randomness Probabilistic polynomials Counting classes are at least as hard as the polynomial-time hierarchy Complexity classes defined by counting quantifiers Relativized circuit complexity Some consequences of non-uniform conditions on uniform classes Robustness of probabilistic computational complexity classes under definitional perturbations --TR --CTR Christian Glaer , Lane A. Hemaspaandra, A moment of perfect clarity II: consequences of sparse sets hard for NP with respect to weak reductions, ACM SIGACT News, v.31 n.4, p.39-51, Dec. 2000 Valentine Kabanets , Jin-Yi Cai, Circuit minimization problem, Proceedings of the thirty-second annual ACM symposium on Theory of computing, p.73-79, May 21-23, 2000, Portland, Oregon, United States Lane A. Hemaspaandra , Mitsunori Ogihara , Gerd Wechsung, Reducing the number of solutions of NP functions, Journal of Computer and System Sciences, v.64 n.2, p.311-328, March 2002 Rahul Santhanam, Circuit lower bounds for Merlin-Arthur classes, Proceedings of the thirty-ninth annual ACM symposium on Theory of computing, June 11-13, 2007, San Diego, California, USA Lane A. Hemaspaandra, SIGACT News complexity theory column 32, ACM SIGACT News, v.32 n.2, June 2001 Jin-Yi Cai , Venkatesan T. Chakaravarthy , Lane A. Hemaspaandra , Mitsunori Ogihara, Competing provers yield improved Karp-Lipton collapse results, Information and Computation, v.198 n.1, p.1-23, April 10, 2005 Piotr Faliszewski , Lane Hemaspaandra, Open questions in the theory of semifeasible computation, ACM SIGACT News, v.37 n.1, March 2006 Lance Fortnow, Beyond NP: the work and legacy of Larry Stockmeyer, Proceedings of the thirty-seventh annual ACM symposium on Theory of computing, May 22-24, 2005, Baltimore, MD, USA Johannes Kbler , Rainer Schuler, Average-case intractability vs. worst-case intractability, Information and Computation, v.190 n.1, p.1-17, April 10, 2004

polynomial-size circuits;lowness;advice classes;randomized computation

298690

Optimal Parallel Algorithms for Finding Proximate Points, with Applications.

AbstractConsider a set P of points in the plane sorted by x-coordinate. A point p in P is said to be a proximate point if there exists a point q on the x-axis such that p is the closest point to q over all points in P. The proximate point problem is to determine all the proximate points in P. Our main contribution is to propose optimal parallel algorithms for solving instances of size n of the proximate points problem. We begin by developing a work-time optimal algorithm running in O(log log n) time and using ${{n \over {\log \log n}}}$ Common-CRCW processors. We then go on to show that this algorithm can be implemented to run in O(log n) time using ${{n \over {\log n}}}$ EREW processors. In addition to being work-time optimal, our EREW algorithm turns out to also be time-optimal. Our second main contribution is to show that the proximate points problem finds interesting, and quite unexpected, applications to digital geometry and image processing. As a first application, we present a work-time optimal parallel algorithm for finding the convex hull of a set of n points in the plane sorted by x-coordinate; this algorithm runs in O(log log n) time using ${{n \over {\log \log n}}}$ Common-CRCW processors. We then show that this algorithm can be implemented to run in O(log n) time using ${{n \over {\log n}}}$ EREW processors. Next, we show that the proximate points algorithms afford us work-time optimal (resp. time-optimal) parallel algorithms for various fundamental digital geometry and image processing problems. Specifically, we show that the Voronoi map, the Euclidean distance map, the maximal empty circles, the largest empty circles, and other related problems involving a binary image of size nn can be solved in O(log log n) time using$${{{n^2} \over {\log \log n}}}$$Common-CRCW processors or in O(log n) time using ${{{n^2} \over {\log n}}}$ EREW processors.

Introduction Consider a parallel algorithm that solves an instance of size n of some problem in T p (n) time using p processors. Traditionally, the main complexity measure for assessing the performance of Work supported in part by NSF grant CCR-9522093, by ONR grant N00014-97-1-0526, and by Grant-in-Aid for Encouragement of Young Scientists (08780265) from Ministry of Education, Science, Sports, and Culture of Japan y Dept. of Electrical and Computer Engineering, Nagoya Institute of Technology, Showa-ku, Nagoya 466, JAPAN, z Department of Computer Science, Old Dominion University, Norfolk, Virginia 23529, USA, [email protected] the algorithm is the amount W (n) of work performed by the algorithm, defined as the product (n). The algorithm is termed work-optimal if W (n) 2 \Theta(T (n)), where T (n) is the running time of the fastest sequential algorithm for the problem. The algorithm is work-time optimal [20] if it is work-optimal and, in addition, its running time T p (n) is best possible among the work-optimal algorithms in that model. Needless to say that one of the challenges of parallel algorithm design is to produce not only work-optimal but, indeed, whenever possible, work-time optimal algorithms. Occasionally, an even stronger complexity metric is being used - the so-called time-optimality. Specifically, an algorithm is time-optimal in a given model, if the problem cannot be solved faster in that model, even if an unbounded number of processors were available. In this paper we assume the Parallel Random Access Machine (PRAM, for short) which consists of synchronous processors, each having access to a common memory. We refer the interested reader to [20] for an excellent discussion of the PRAM model. Let P be a set of points in the plane sorted by x-coordinate. A point p is a proximate point of P if there exists a point on the x-axis closer to p than to any other point in P . The proximate points problem asks to determine all proximate points in P . Clearly, the proximate points problem can be solved, using an algorithm for finding the Voronoi diagram. However, as argued in [16], the computation of the Voronoi diagram log n) time even if the n points are sorted by x-coordinate. Thus, this naive approach does not yield an optimal solution to the proximate points problem. Recently, Breu et al. [6] proposed a linear-time algorithm for the proximate points problem. In spite of its optimality, the algorithm of Breu et al. [6] relies in crucial ways on stack operations, notoriously hard to parallelize. Our first main contribution is to propose parallel algorithms for solving instances of size n of the proximate points problem. Specifically, we first exhibit an algorithm running in O(log log n) time using n log log n Common-CRCW processors. We then go on to show that this algorithm can be implemented to run in O(log n) time using n log n EREW processors. Our Common-CRCW algorithm is work-time optimal; the EREW algorithm turns out to also be time-optimal. We establish the work-time optimality of our Common-CRCW algorithm by a reduction from the minimum finding problem; the time-optimality of our EREW algorithm follows by a reduction from the OR problem. Our second main contribution is to show that the proximate points problem has interesting, and quite unexpected, applications to digital geometry and image processing. To begin, we present a work-time optimal parallel algorithm for computing the convex hull of a set of n points in the plane sorted by x-coordinate. This algorithm runs in O(log log n) time using n log log n Common- CRCW processors or in O(log n) time using n log n EREW processors. We show that this algorithm is work-time optimal in the CRCW model and, in addition, time-optimal in the EREW. Numerous parallel algorithms have been proposed for computing the convex hull of sorted points in the plane [4, 7, 12, 13, 21]. Recently, Chen [7] presented an O(log n)-time algorithm using n log n processors. Chen et al. [12] presented work-optimal algorithms running in O(log n)-time algorithm and using n log n EREW processors, and in an O(log log n)-time algorithm using n log log n Common-CRCW processors. Quite recently, Berkman et al. [3] presented an O(log log n)-time algorithm using n log log n Common-CRCW processors. Our algorithm features the same performance as those in [3, 7, 12]. However, our algorithm is much simpler and more intuitive. Further, to the best of our knowledge, the work-time optimality of the CRCW version and the time-optimality of the EREW version algorithm has not been solved yet. Given a binary image the Voronoi map assigns to each pixel in the image the position of the nearest black pixel. The Euclidean distance map assigns to each pixel the Euclidean distance to the nearest black pixel. An empty circle of the image is a circle whose interior contains only white pixels. A maximal empty circle is an empty circle contained in no other empty circle. A largest empty circle is an empty circle of the largest radius. We refer the reader to Figure 1 for an illustration. The largest square, diamond, n-gon, etc. are defined similarly. These computations are known to have numerous applications ranging from clustering and shape analysis [2, 17] to handoff management in cellular systems [26] to image compression, decomposition, and reconstruction [5, 23, 27, 28, 31]. As further applications, we propose algorithms for computing the Voronoi map, the Euclidean distance map, the maximal empty circles, and the largest empty circles of a binary image of size n \Theta n. We begin by presenting a work-time optimal algorithm that computes the Voronoi map and the Euclidean distance map of a binary image of size n \Theta n in O(log log n) time using n 2 log log n Common-CRCW processors or in O(log n) time using n 2 log n EREW processors. We also show that the distance map for various metrics including the well known L k metrics, (k - 1), can also be computed in the same manner. We then go on to show that all the maximal empty circles and a largest empty circle of an n \Theta n binary image can be found in O(log log n) time using n 2 log log n Common-CRCW processors or in O(log n) time using n 2 log n EREW processors. As it turns out, with minimal changes, this algorithm is applicable to various other kinds of empty figures including squares, diamonds, n-gon etc. Recently, Chen et al. [8, 11] and Breu et al. [6] presented O(n 2 )-time sequential algorithms for computing the Euclidean distance map. Roughly at the same time, Hirata [19] presented a simpler sequential algorithm to compute the distance map for various distance metrics including Euclidean, 4-neighbor, 8-neighbor, chamfer, and octagonal. A number of parallel algorithms for computing the Euclidean distance map have been developed for various parallel models [1,9,10,14]. In particular, the following results have been reported in the recent literature. Lee et al. [22] Figure 1: Illustrating the Euclidean distance map and the largest empty circle presented an O(log 2 n)-time algorithm using n 2 EREW processors. Pavel and Akl [24] presented an algorithm running in O(log n) time and using n 2 EREW processors. Clearly, these two algorithms are not work-optimal. Chen [8] presented a work-optimal O( n 2 )-time algorithm using p, (p log p - n), EREW processors. This yields an O(n log n)-time algorithm using n log n EREW processors. Fujiwara et al. [18] presented a work-optimal algorithm running in O(log n) time and using n 2 log n EREW processors and in O( log n log log n ) time using n 2 log log n log n Common-CRCW processors. Although Fujiwara et al. [18] claim that their algorithm is applicable to various distance maps, a closer analysis reveals that it only applies to a few distance metrics. The main problem seems to be that their algorithm uses a geometric transform that depends in a crucial way on properties of the Euclidean distance and, therefore, does not seem to generalize. As we see it, our Euclidean distance map algorithm has three major advantages over Fujiwara's algorithm. First, the performance of our algorithm for the CRCW is superior; second, our algorithm applies to a large array of distance finally, our algorithm is much simpler and more intuitive. The remainder of this paper is organized as follows: Section 2 introduces the proximate points problem for the Euclidean distance metric and discusses a number of technicalities that will be crucial ingredients in our subsequent algorithms. Section 3 presents our parallel algorithms for the Common-CRCW and the EREW. Section 4 proves that these algorithms are work-time, respectively time-optimal. Section 5 presents a work-time optimal parallel algorithm for computing the convex hull of sorted points in the plane. Section 6 uses the proximate points algorithm to computing the Voronoi map, the Euclidean distance map, the maximal empty circles, and the largest empty circles of a binary image. Section 7 offers concluding remarks and open problems. Finally, the Appendix discusses other distance metrics to which the algorithms presented in Section 3 apply. 2 The proximate points problem: a first look In this section we introduce the proximate points problem along with a number of geometric results that will lay the foundation of our subsequent algorithms. Throughout, we assume that a point p is represented by its Cartesian coordinates (x(p); y(p)). As usual, we denote the Euclidean distance between the planar points p and q by d(p; Consider a collection of n points sorted by x-coordinate, that is, such that We assume, without loss of generality that all the points in P have distinct x-coordinates and that all of them lie above the x-axis. The reader should have no difficulty confirming that these assumptions are made for convenience only and do not impact the complexity of our algorithms. Recall that for every point p i of P the locus of all the points in the plane that are closer to than to any other point in P is referred to as the Voronoi polygon associated with p i and is denoted by V (i). The collection of all the Voronoi polygons of points in P partitions the plane into the Voronoi diagram of P (see [25] p. 204). Let I i , (1 - i - n), be the locus of all the points q on the x-axis for which d(q; In other words, q 2 I i if and only if q belongs to the intersection of the x-axis with V (i), as illustrated in Figure 2. In turn, this implies that I i must be an interval on the x-axis and that some of the intervals I i , may be empty. A point p i of P is termed a proximate point whenever the interval I i is nonempty. Thus, the Voronoi diagram of P partitions the x-axis into proximate intervals. Since the points of are sorted by x-coordinate, the corresponding proximate intervals are ordered, left to right, as point q on the x-axis is said to be a boundary point between p i and p j if q is equidistant to p i and p j , that is, d(p It should be clear that p is a boundary point between proximate points p i and p j if and only if the q is the intersection of the (closed) intervals I i and I j . To summarize the previous discussion we state the following result. Proposition 2.1 The following statements are satisfied: ffl Each I i is an interval on the x-axis; ffl The intervals I 1 ; I lie on x-axis in this order, that is, for any non-empty I i and I j lies to the left of I I 1 I 2 I 4 I 6 I 7 Figure 2: Illustrating proximate intervals I 1 I 2 I 3 I 4 I 0 Figure 3: Illustrating the addition of p to g. ffl If the non-empty proximate intervals I i and I j are adjacent, then the boundary point between Referring again to Figure 2, among the seven points, five points are proximate points, while the others are not. Note that the leftmost point p 1 and the rightmost point p n are always proximate points. Given three points we say that dominated by p i and p k whenever fails to be a proximate point of the set consisting of these three points. Clearly, p j is dominated by p i and p k if the boundary of p i and p j is to the right of that of p j and p k . Since the boundary of any two points can be computed in O(1) time, the task of deciding for every triple (p whether p j is dominated by p i and p k takes O(1) time using a single processor. Consider a collection of points in the plane sorted by x-coordinate, and a point p to the right of P , that is, such that x(p 1 x(p). We are interested in updating the proximate intervals of P to reflect the addition of p to P as illustrated in Figure 3. We assume, without loss of generality, that all points in P are proximate points and let I n be the corresponding proximate intervals. Further, let I 0 p be the up-dated proximate intervals of P [ fpg. Let p i be a point such that I 0 i and I 0 are adjacent. By (iii) in Proposition 2.1, the boundary point between p i and p separates I 0 i and I 0 . As a consequence, (ii) implies that all the proximate intervals I 0 n must be empty. Furthermore, the addition of p to P does not affect any of the proximate intervals I j , 1 In other words, for all 1 I 0 are empty, the points p are dominated by p i and p. Thus, every point n), is dominated by otherwise, the boundary between would be to the left of that of that between p j and p. This would imply that the non-empty interval between these two boundaries corresponds to I 0 j , a contradiction. To summarize, we have the following result. Lemma 2.2 There exists a unique point p i of P such that: ffl The only proximate points of P [ fpg are ffl For the point p j is not dominated by I 0 ffl For dominated by and the interval I 0 j is empty. i and I 0 are consecutive on the x-axis and are separated by the boundary point between and p, be a collection of proximate points sorted by x-coordinate and let p be a point to the left of P , that is, such that x(p) ! x(p 1 reference we now take note of the following companion result to Lemma 2.2. The proof is identical and, thus, omitted. Lemma 2.3 There exists a unique point p i of P such that: ffl The only proximate points of P [ fpg are ffl For not dominated by p and p j+1 . Moreover, for I 0 ffl For the point p j is dominated by p and p j+1 and the interval I 0 j is empty. p and I 0 are consecutive on the x-axis and are separated by the boundary point between p and The unique point p i whose existence is guaranteed by Lemma 2.2 is termed the contact point between P and p. The second statement of Lemma 2.2 suggests that the task of determining the unique contact point between P and a point p to the right or left of P reduces, essentially, to binary search. Now, suppose that the set into two subsets g. We are interested in updating the proximate intervals in the process or merging PL and PR . For this purpose, let I 2n be the proximate intervals of PL and PR , respectively. We as- sume, without loss of generality, that all these proximate intervals are nonempty. Let I 0 be the proximate intervals of . We are now in a position to state and prove the next result which turns out to be a key ingredient in our algorithms. Lemma 2.4 There exists a unique pair of proximate points PR such that ffl The only proximate points in PL [ PR are are empty, and I 0 ffl The proximity intervals I 0 i and I 0 are consecutive and are separated by the boundary point between Proof. Let i be the smallest subscript for which p i 2 PL is the contact point between PL and a point in PR . Similarly, let j be the largest subscript for which the point PR is the contact point between PR and some point in PL . Clearly, no point in PL to the left of p i can be a proximate point of P . Likewise, no point in PR to the left of p j can be a proximate point of P . Finally, by Lemma 2.2 every point in PL to the left of p i must be a proximate point of P . Similarly, by Lemma 2.3 every point in PR to the right of p i must be a proximate point of P , and the proof of the lemma is complete. The points p i and p j whose existence is guaranteed by Theorem 2.4 are termed the contact points between PL and PR . We refer the reader to Figure 4 for an illustration. Here, the contact points between PL and PR are p 4 and p 8 . Next, we discuss a geometric property that enables the computation of the contact points p i and p j between PL and PR . For each point p k of PL , let q k denote the contact point between p k and PR as specified by Lemma 2.3. We have the following result. Lemma 2.5 The point p k is not dominated by p k\Gamma1 and q k if 2 - k - i, and dominated otherwise. I 1 I 2 I 3 I 4 I 5 I 6 I 7 I 8 I 9 I 10 I 0 9 I 0 Figure 4: Illustrating the contact points between two sets of points Proof. If dominated by p k\Gamma1 and q k , then I 0 k must be empty. Thus, Lemma 2.4 guarantees that p k , (2 - k - i), is not dominated by p k\Gamma1 and q k . Suppose that p k , (i is not dominated by p k\Gamma1 and q k . Then, the boundary point between p k and q k is to the right of that between p k\Gamma1 and p k . Thus, the non-empty interval between these two boundaries corresponds to I 0 k , a contradiction. Therefore, p k , (i n), is dominated by p k\Gamma1 and q k , completing the proof. Lemma 2.5 suggests a simple, binary search-like, approach to finding the contact points p i and between two sets PL and PR . In fact, using a similar idea, Breu et al. [6] proposed a sequential algorithm that computes the proximate points of an n-point planar set in O(n) time. The algorithm in [6] uses a stack to store the proximate points found and, consequently, seems very hard to parallelize. 3 Parallel algorithms for the proximate points problem We begin by discussing a parallel algorithm for solving the proximate points problem on the Common-CRCW. The algorithm will then be converted to run on the EREW. We rely, in part, on the solution to the well-known LEFTMOST-ONE problem: given a sequence b 1 determine the smallest i, (1 - i - n), such that b Lemma 3.1 [20] An instance of size n of the LEFTMOST-ONE problem can be solved in O(1) time using n Common-CRCW processors. Consider a set of points such that x(p 1 To capture the neighboring proximate points of each point use three indices c i , l i and r i defined as follows: non-proximate points proximate points Figure 5: Illustrating indices l i , c i , and r i for a point p i 1. is an proximate pointg; 2. l is an proximate pointg; 3. r is an proximate pointg. We refer the reader to Figure 5 for an illustration. Note that we must have l there is no proximate point p j such that l i then c Next, we are interested in finding the contact point between the set and a new point p with x(p n We assume that for every i, (1 - i - n), c i , l i , and r i are available, and that m, (m - n), processors are at our disposal. The algorithm is essentially performing m-ary search using Lemma 2.2. Algorithm Find-Contact-Point(P; p) Extract a sample S(P ) of size m consisting of the points p c in P . For every k, (k - 0), check whether the point p c k n is dominated by p l k n and p, and whether is dominated by p c k n and p. If p c k n is not dominated but p r k n is dominated, then is the desired contact point. such that the point p r k n is not dominated by p c k n and p, and p c (k+1) n is dominated by p l (k+1) n and p. Step 3 Execute recursively this algorithm for the set of points P g to find the contact point. 1, the set P 0 contains at most l (k+1) n points. Hence, the depth of the recursion is O( log n log m ). Notice, further, that algorithm Find-Contact-Point does not perform concurrent reading or writing. Thus, we have the following result. Lemma 3.2 Given a set of n points in the plane sorted by x-coordinate and a point p, with x(p n the task of finding the contact point between P and p can be performed in O( log n log using m EREW processors. Next, consider two sets of points in the plane such that x(p 1 Assume that for every i the indices c i , l i and r i are given and that m processors are available to us. The following algorithm finds the contact points of PL and PR by m-ary search using Lemma 2.5. Algorithm Find-Contact-Points-Between-Sets(PL ; PR ) sample points S(PL from PL . By using the algorithm Find-Contact-Point and p m of the processors available each, determine for each sample point 1), the corresponding contact point q c k n in PR . Step 2 For each k, (0 - k - check whether the point p c k n is dominated by p l k n and q c k n , and whether the point p r k n is dominated p c k n and q c k n . If p c k n is not dominated, yet p r k n is, output and q c k n as the desired contact points. Step 3 Find k such that the point p r k n is not dominated by p c k n and q c k n is dominated by p l (k+1) n and q c (k+1) n Step 4 Execute recursively algorithm Find-Contact-Points-Between-Sets for the sets P 0 and PR and return the desired contact points. It is not hard to see that algorithm Find-Contact-Points-Between-Sets involves concurrent reads (because several processors may access a point concurrently), but does not involve concurrent operations. By Lemma 3.2, Step 1 can be takes O( log n log m ) time on the CREW model. Steps 2 and 3 run, clearly, in O(1) time. Since P 0 L contains at most n the depth of recursion is O( log n log m ). Thus, altogether, algorithm Find-Contact-Points-Between-Sets runs in O( log 2 n using m CREW processors. Lemma 3.3 Given the sets of points in the plane such that x(p 1 the task of finding the contact points between PL and PR can be performed in O( log 2 n using m CREW processors. Next, we are interested in designing an algorithm to compute the proximate points of a set P on n points in the plane sorted by x-coordinate in O(log log n) time on the Common-CRCW. We assume that n processors are available to us. We begin by determining for every i, the indices c i , l i , and r i . With this information available, all that remains to be done is to retain all the points p i for which c i. The details follow. Algorithm Find-Proximate-Points(P) Partition the set P into n 1=3 subsets such that for every k, (0 - k - g. For every point p i in P k , (0 - k - n 1=3 \Gamma 1), determine the indices c i , l i , and r i local to P k . Compute the contact points of each pair of sets P i and using n 1=3 of the processors available. Let q i;j 2 P i denote the contact point between P i and P j . Step 3 For every P i , find the rightmost contact point p rc i among all the points q i;j with and find the leftmost contact point p lc i over all points q i;j with j ? i. Clearly, x(p rc i ig. Step 4 For each set P i , the proximate points lying between rc i and lc i (inclusive) are proximate points of P . Update each c It is clear that Step 2 can be performed in O( log n 2=3 log runs in O(1) time using Lemma 3.1. At this moment, the reader may wonder how the updating of the indices can be performed efficiently. In fact, as it turns out, this update can be done in O(1) time. Since the task of updating l i and r i is, essentially, the same as that of updating c i , we will only focus on c i . In each P i , for all the points p j , (rc the value of c j is not changed. For all the points p j with lc i ! j, the value of c j must be changed to lc i . For all points p j with the value of c j is changed to lc has an proximate point. However, if P contains no proximate points, we have to find the nearest subset that contains a proximate point. To do this, first check whether each P i has a proximate point using n 2=3 processors each. Thus, totally, processors are used for this task. Next, using Lemma 3.1 we determine k such that contains a proximate pointg for each P i . Since P has n 1=3 groups, this task can be done in O(1) time and n 1=3 processors each and, totally, n 1=3 \Delta n processors are used. Thus, Step 4 can be done in O(1) time using n processors. Let TCRCW (n) be the running time of this algorithm. To find the recurrence describing the worst case running time of algorithm Find-Proximate-Points, we note that Step 1 executes recursively this algorithm for n 2=3 points, while Steps 2, 3, and 4 run in O(1) time. Thus, we have confirming that TCRCW (n) 2 O(log log n). Thus, we have: Lemma 3.4 An instance of size n of the proximate points problem can be solved in O(log log n) time using n Common-CRCW processors. Next, we show that the number of processors can be reduced by a factor of log log n without increasing the running time. The idea is as follows: begin by partitioning the set P into n log log n subsets log log n each of size log log n. Next, using algorithm Sequential-Proximate-Points find the proximate points within each subset in O(log log n) sequential time and, in the process, remove from P all the points that are not proximate points. For every i, (1 - i - n log log n ), let proximate points in the set P i . At this moment, execute algorithm Find-Proximate-Point on P 1 log log n . Since n processors are required in order to update the indices c i , l i , and r i in O(1), we will proceed slightly differently. The idea is the following: while executing the algorithm, some of the (currently) proximate points will cease to be proximate points. To maintain this information efficiently, we use ranges log log n log log n such that for each P i , fp i;L are the current proximate points. While executing the algorithm, P i may contain no proximate points. To find the neighboring proximate points, we use the pointers L 0 log log n and log log n such that and the set P j contains a proximate pointg, and the set P j contains a proximate pointg. By using this strategy, we can find the contact point between a point and P in O( log n log using processors as discussed in Lemma 3.2. Thus, the contact points between two subsets can be found in the same manner as in Lemma 3.3. Finally, the algorithm for Lemma 3.1 can update i in Step 4 in O(1) time by using O( n log log n ) processors. To summarize, we have the following result. Theorem 3.5 An instance of size n of the proximate points problem can be solved in O(log log n) time using n log log n Common-CRCW processors. We close this section by pointing out that algorithm Find-Proximate-Points can be implemented efficiently on the EREW. For this purpose, we rely, in part, on the following well known result [20]. Lemma 3.6 A single step execution of the m-processor CRCW can be simulated by an m-processor EREW in O(log m) time. By Lemma 3.6, Steps 2, 3, and 4 of the algorithm can be performed in O(log n) time using processors, as the CRCW performs these steps in O(1) time using n processors. Let TEREW (n) be the worst-case running time on the EREW. Then, the recurrence describing the confirming that T (n) 2 O(log n). Consequently, we have: Lemma 3.7 An instance of size n of the proximate points problem can be solved in O(log n) time using n EREW processors. Using, essentially, the same idea as for the Common-CRCW, we can reduce the number of processors by a factor of log n without increasing the computing time. Specifically, in case of the EREW, the n points are partitioned into n log n subsets each of size log n. Thus, we have Theorem 3.8 An instance of size n of the proximate points problem can be solved in O(log n) time using n log n EREW processors. 4 Lower Bounds The main goal of this section is to show that the running time of the Common-CRCW algorithm for the proximate points problem developed in Section 3 cannot be improved while retaining work- optimality. This, in effect, will prove that our Common-CRCW algorithm is work-time optimal. We then show that our EREW algorithm is time-optimal. The work-optimality of both algorithms is obvious; in order to solve the proximate points problem every point must be accessed at least once. n) work is required of any algorithm solving the problem. Our lower bound arguments rely, in part, on the following fundamental result of Valiant [30]. Lemma 4.1 The task of finding the minimum (maximum) of n real numbers log n) time on the CRCW provided that n log O(1) n processors are available. We now show that the lower bound of Lemma 4.1 holds even if all the item are non-negative. Lemma 4.2 The task of finding the minimum (maximum) of n non-negative (non-positive) real numbers requires log n) time on the CRCW provided that n log O(1) n processors are available. Proof. Assume that the minimum (maximum) of n non-negative numbers can be computed in o(log log n) time using n log O(1) n CRCW processors. With this assumption, we can find the minimum of n real numbers in o(log log n) time as follows: first, in O(1) time, check whether there are negative numbers in the input. If not, the minimum of input items can be computed in o(log log n) time. If negative numbers exist, replace every non-positive number by 0 and find the maximum of their absolute values in the resulting sequence in o(log log n) time. The maximum thus computed corresponds to the minimum of the original input. Thus, the minimum of n real numbers can be computed in o(log log n) time, contradicting Lemma 4.1. Further, we rely on the following classic result of Cook et al. [15]. Lemma 4.3 The task of finding the minimum (maximum) of n real numbers on the CREW (therefore, also on the EREW) even if infinitely many processors are available. We shall reduce the task of finding the minimum of a collection A of n non-negative a 1 ; a to the proximate points problem. Our plan is to show that an instance of size n of the problem of finding the minimum of a collection of non-negative numbers can be converted, in O(1) time, to an instance of size 2n of the the proximate points problem involving sorted points in the plane. For this purpose, let be a set of arbitrary non-negative real numbers that are input to the minimum problem. We construct a set of points in the plane by setting for every i, (1 Notice that this construction guarantees that the points in P are sorted by x-coordinate and that for every i, (1 - i - n), the distance between the point p i and the origin is exactly Intuitively, our construction places the 2n points circles centered at the origin. More precisely, for every i, (1 - i - n), the points p i and p n+i are placed on such a circle C i with radius . It is very important to note that the construction above can be carried out in O(1) time using n EREW processors. In our subsequent arguments, we find it convenient to rely on the next technical result. Lemma 4.4 Both p i and p i+n are proximate points if and only if a i is the minimum of A. Proof. Let a i be the minimum of A and refer to Figure 6. Clearly, C i is the circle of smallest radius containing p i and p i+n , while all the other points lie outside C i . Hence, p i and p i+n are the closest points of P from the origin. Thus, the boundary between lies to the left of the origin: were this not true, p j would be closer to the origin than p i+n . The following simple facts are proved in essentially the same way. a i a i+n I i+n I i O Figure Illustrating P for Lemma 4.4 1. The boundary point between p j and p i+n lies to the left of the origin if 1 and to the right if i 2. The boundary point between p j and p i lies to the left of the origin if 1 and to the right if for each point n), the boundary between p j and p i lies to the left of that between p j and p n+i , p j is not proximate point. Thus, for j 6= i, either p j or p j+n fails to be a proximate point. Further, for the point p i the boundary with lies to the left of the origin, and that with n) lies to the right of the origin (or is the origin itself). Thus, p i is a proximate point. The fact that p i+n is a proximate point follows by a mirror argument. This completes the proof. Lemma 4.4 guarantees that the minimum of A can be determined in O(1) time once the proximate points of P are known. Now, Lemma 4.1 implies the following important result. Theorem 4.5 Any algorithm that solves an instance of size n of the proximate points problem on the CRCW must take \Omega\Gammake/ log n) time in the worst case, provided that n log O(1) n processors are available. Using exactly the same construction, in combination with Lemma 4.3 we obtain the following lower bound for the CREW. Theorem 4.6 Any algorithm that solves an instance of size n of the proximate points problem on the CREW (also on the EREW) must take \Omega\Gammake/ n) time, even if an infinite number of processors are available. Notice that the EREW algorithm for the proximate points problem presented in Section 3 running in O(log n) time using n log n processors features the same work and time performance on the CREW-PRAM. By Theorem 4.6 the corresponding CREW algorithm is also time-optimal. It is straightforward to extend the previous arguments to handle the case of the L k metric. Specifically, in this case, for every i, (1 - i - n), the points allow us to find the minimum of A. Thus, Theorems 4.5 and 4.6 provide lower bounds for solving the proximate points problem for the distance metric L k . 5 Computing the convex hull The main goal of this section is to show that the proximate points algorithms developed in Section 3 yield a work-time optimal (resp. time-optimal) algorithm for computing the convex hull of a set of points in the plane sorted by x-coordinate. We begin by discussing the details of this algorithm. In the second subsection we establish its work-time (resp. time) optimality. 5.1 The convex hull algorithm be a set of n points in the plane with x(p 1 line segment partitions the convex hull of P into the lower hull, lying below the segment, and the upper hull, lying above it. We focus on the computation of the lower hull only, the computation of the upper hull being similar. For a sequence a 1 , a of items, the prefix maxima is the sequence a 1 , maxfa 1 ; a g. For later reference, we state the following result [20, 29]. Lemma 5.1 The task of computing the prefix maxima (prefix minima) of an n-item sequence can be performed in O(log n) time using n log n EREW processors or in O(log log n) time using n log log n Common-CRCW processors. be a set of n points in the plane sorted by x-coordinate as x(p 1 We define a set let of n points by setting for every i, (1 - i - n), q It is important to note that the points in Q Figure 7: Illustrating the proof of Lemma 5.2 are also sorted by x-coordinate. The following surprising result captures the relationship between the sets P and Q we just defined. Lemma 5.2 For every j, (1 - j - n), p j is an extreme point of the lower hull of P if and only if q j is a proximate point of Q. Proof. If is an extreme point of P and q j is a proximate point of Q. Thus, the lemma is correct for Now consider an arbitrary j in the range be arbitrary subscripts such that 1 - be the boundaries between q i and q j , and between q j and q k , respectively, and refer to Figure 7. Clearly, Thus, we have Similarly, we obtain It is easy to see that the point q j is not dominated by q i and q k if and only if x(b Notice that the slopes of the line segments are 2x(b i ) and 2x(b k ), respectively. Thus, the point p j lies below the segment p i p k if and only if 2x(b Consequently, the point lies below the segment p i p k if and only if the point q j is not dominated by q i and q k . In other words, the point p j is an extreme point of the lower hull of P if and only if q j is a proximate point of Q. Lemma 5.2 suggests the following algorithm for determining the extreme points of the lower hull of g. Algorithm Find-Lower-Hull(P ) Construct the set by setting for every i, (1 - i - n), Determine the proximate points of Q and report p i as an extreme point of the lower hull of P whenever q i is a proximate point of Q. The preprocessing in Step 0 amounts to translating the set P vertically in such a way that for every This affine transformation does not affect the convex hull of P . The correctness of this simple algorithm follows directly from Lemma 5.2. To argue for the running time, we note that by Lemma 5.1 Step 0 takes O(log log n) time and optimal work on the Common-CRCW or O(log n) time and optimal work on the EREW. Step 1 runs in O(1) time using optimal work on either the Common-CRCW or the EREW. By Theorems 3.5 and 3.8, Step 2 takes O(log log n) time and optimal work on the Common-CRCW or O(log n) time and optimal work on the EREW. Thus, we have proved the following result. Theorem 5.3 The task of computing the convex hull of a set of n points sorted by x-coordinate can be performed in O(log log n) time using n log log n Common-CRCW processors or in O(log n) time using n log n EREW processors. 5.2 The optimality of the convex hull algorithm The main goal of this subsection is to show that the convex hull algorithm described in the previous subsection is work-time optimal on the Common-CRCW and, in addition, time-optimal on the CREW and EREW. Clearly, every point must be read at least once to solve the proximate points problem. Thus, O(n)-time is required to solve the problem, and our convex hull algorithms (Common-CRCW or are work-optimal. Next, we show that given a set of n non-negative integers their maximum can be determined by using any algorithm for computing the convex hull of a set of sorted points in the plane. For this purpose, we exhibit an O(1)-time reduction of the maximum problem to the convex hull problem. The proof technique is similar to the one employed for the proximate points problem. With A given construct a set of points in the plane by setting for every i, is a set of point in the plane sorted by x-coordinate. The following result relates the sets A and P . Lemma 5.4 The item a i is the maximum of A if and only if both p i and p i+n are points on the upper hull of P . Proof. Let a i be the maximum of A. By construction, both p i and p i+n are points of the upper hull of P . Further, none of the points p can belong to the upper hull of P . Thus, there exist no subscript j, (j 6= i), for which both p j and p n+i belong to the upper hull of P . This completes the proof. Consequently, to find the maximum of A all we need do is to find an index i such that both p i and p n+i are points of the upper hull. Therefore, the problem of finding the upper hull of 2n sorted points in the plane is at least as hard as the problem of finding the maximum of n non-negative numbers. Thus, we have the following important result. Theorem 5.5 The task of finding the convex hull of n points log n) time on the CRCW, provided that n log O(1) n processors are available. Similarly, we have the following companion result. Theorem 5.6 The task of finding the convex hull of n points n) time on the CREW, even if infinitely many processors are available. By Theorems 5.5 and 5.6 the convex hull algorithms developed in the previous section are work-time optimal. In addition, the EREW algorithm is both work-time and time-optimal. 6 Applications to image processing A binary image I of size n \Theta n is maintained in an array b i;j , (1 - n). It is customary to refer to pixel (i; j) as black if b The rows of the image will be numbered bottom up starting from 1. Likewise, the columns will be numbered left to right, with column 1 being the leftmost. In this notation pixel b 1;1 is in the south-west corner of the image. The Voronoi map associates with every pixel in I the closest black pixel to it (in the Euclidean metric). More formally, the Voronoi map of I is a function I such that for every (i; j), only if where is the Euclidean distance between pixels (i; The Euclidean distance map of image I associates with every pixel in I the Euclidean distance to the closest black pixel. Formally, the Euclidean distance map is a function R such that for every (i; j), (1 - In our subsequent arguments we find it convenient to rely on the solution to the NEAREST-ONE problem: given a sequence of 0's and 1's, determine the closest 1 to every item in A. As a direct corollary of Lemma 5.1 we have Lemma 6.1 An instance of size n of the NEAREST-ONE problem can be solved in O(log log n) time using n log log n Common-CRCW processors or in O(log n) time using n log n EREW processors. We assume a binary image I of size n \Theta n as discussed above and the availability of n 2 processors, where T log n) for the Common-CRCW and T n) for the EREW. We now outline the basic idea of our algorithm for computing the Voronoi map and the Euclidean distance map of image I. We begin by determining, for every pixel in row j, (1 - j - n), the nearest black pixel, if any, in the same column of the subimage of I. More precisely, with every pixel (i; j) we associate the value Next, we construct an instance of the proximate points problem for every row j, (1 - j - n), in the image I involving the set P j of points in the plane defined as P Having solved, in parallel, all these instances of the proximate points problem, we determine, for every proximate point p i;j in P j its corresponding proximity interval I i . With j fixed, we determine for every pixel (i; (that we perceive as a point on the x-axis) the identity of the proximity interval to which it belongs. This allows each pixel (i; j) to determine the identity of the nearest pixel to it. The same task is executed for all rows in parallel, to determine for every pixel (i; in row j the nearest black pixel. The details are spelled out in the following algorithm. Algorithm Voronoi-and-Euclidean-Distance-Map(I) Step 1 For each pixel (i; j), compute the distances ng to the nearest black pixel in the same column as (i; j) in the subimage of I. Step 2 For every j, (1 ng. Compute the proximate points E(P j ) of P j . Step 3 For every point p in E(P j ) determine its proximity interval of P j . Step 4 For every i, (1 - i - n), determine the proximate intervals of P j to which the point (i; (corresponding to pixel (i; j)) belongs. The correctness of this algorithm being easy to see we turn to the complexity. Step 1 can be performed in O(T (n)) time using the processors available by using Lemma 6.1. Theorem 3.5 and 3.8 guarantee that Step 2 takes O(T using n processors. By Lemma 6.1, Steps 3 and 4 can be performed in the same complexity. Thus, we have the following important result. Theorem 6.2 The task of computing the Voronoi map and the Euclidean distance map of a binary image of size n \Theta n can be performed in O(log log n) time using n 2 log log n Common-CRCW processors or in O(log n) time using n 2 log n EREW processors. Recall that an empty circle in the image I is a circle filled with white pixels. The task of computing the largest empty circles in an image is a recurring theme in pattern recognition, robotics, and digital geometry [17]. An empty circle is said to be maximal if it is contained in no other empty circle. An empty circle is said to me maximum if its radius is as large as possible. It is clear that a maximum empty circle is also a maximal, but not conversely. We now turn to the task of determining all maximal (resp. maximum) empty circles in an input image I. Algorithm All-Maximal-Empty-Circles(I) Compute the Euclidean distance map m of I. Step 2 For each pixel (i; j), (1 - I compute the smallest distance u jg to the border of the image. Then, compute r which is the largest radius of every empty circle centered at the pixel (i; j). Step 3 For each pixel (i; check whether there exists a neighboring pixel (i 1), such that the circle with radius r i;j and origin (i; j) is included by the circle with radius r i 0 ;j 0 and origin (i no such circle exists, label the circle of radius r i;j centered at as a maximal empty circle. g. Every pixel (i; in I for which r its empty circle as the largest empty circle of I. Clearly, all the steps of this simple algorithm can be performed in O(log log n) time using n 2 log log n Common-CRCW processors or in O(log n) time using n 2 log n EREW processors. Thus we have Corollary 6.3 The task of labeling all the maximal empty circles and of reporting a maximum empty circle of a binary image of size n \Theta n can be performed in O(log log n) time using n 2 log log n Common-CRCW or in O(log n) time using EREW n 2 log n processors. Conclusions Our first main contribution is to propose optimal parallel algorithms for solving instances of size n of the proximate points problem. Our first algorithm runs in O(log log n) time and uses n log log n Common-CRCW processors. This algorithm can, in fact, be implemented to run in O(log n) time using n log n EREW processors. The Common-CRCW algorithm is work-time optimal; the EREW algorithm is, in addition, time-optimal. out to also be time-optimal. Our second main contribution is to show that the proximate points problem finds interesting, and quite unexpected, applications to digital geometry and image processing. As a first application we presented a work-time optimal parallel algorithm for finding the convex hull of a set of n points in the plane sorted by x-coordinate; this algorithm has the same complexity as the proximate points algorithm. Next, we showed that the proximate points algorithms afford us work-time optimal (resp. time-optimal) parallel algorithms for various fundamental digital geometry and image processing problems. Specifically, we show that the Voronoi map, the Euclidean distance map, the maximal empty circles, the largest empty circles, and other related problems. Further, we have proved the work-time, respectively, the time optimality of our proximate points and convex hull algorithms. However, for the image processing problems discussed, it is not known whether the algorithms developed are optimal. We conjecture that, for these problems,\Omega\Gammaobl log n) is a time lower bound on the CRCW, provided that the algorithms are work-time optimal. For the CREW and EREW, the logical-OR problem can be reduced to these image processing problems quite easily. Therefore,\Omega\Gammaher n) is a time lower bound for both the CREW and the EREW. --R Euclidean distance transform on polymorphic processor array. Computer Vision. A fast parallel algorithm for finding the convex hull of a sorted point set. Centres of maximal discs in the 5-7-11 distance transform Linear time Euclidean distance transform algorithms. Efficient geometric algorithms on the EREW PRAM. Optimal algorithm for complete Euclidean distance transform. Designing systolic architectures for complete Euclidean distance trans- form An efficient algorithm for complete Euclidean distance transform on mesh-connected SIMD A fast algorithm for Euclidean distance maps of a 2-d binary image Optimal parallel algorithms for computing convex hulls. A parallel method for the prefix convex hulls problem. SIMD hypercube algorithm for complete Euclidean distance transform. Upper and lower time bounds for parallel random access machines without simultaneous writes. On computing Voronoi diagrams for sorted point sets. Pattern Classification and Scene Analysis An optimal parallel algorithm for the Euclidean distance maps. A unified linear-time algorithm for computing distance maps An Introduction to Parallel Algorithms. Efficient parallel geometric algorithms on a mesh of trees. Parallel computation of exact Euclidean distance transform. Modified distance transform with raster scanning value propagation. Efficient algorithms for the Euclidean distance transform. Computational Geometry: An Introduction. A skeletonization algorithm by maxima tracking on Euclidean distance transform. Finding the maximum Parallelism in comparison problem. On the generation of skeletons from discrete Euclidean distance maps. --TR --CTR Ling Chen , Yi Pan , Xiao-hua Xu, Scalable and Efficient Parallel Algorithms for Euclidean Distance Transform on the LARPBS Model, IEEE Transactions on Parallel and Distributed Systems, v.15 n.11, p.975-982, November 2004 Amitava Datta , Subbiah Soundaralakshmi, Fast and scalable algorithms for the Euclidean distance transform on a linear array with a reconfigurable pipelined bus system, Journal of Parallel and Distributed Computing, v.64 n.3, p.360-369, March 2004

parallel algorithms;proximate points;digital geometry;convex hulls;pattern recognition;largest empty circles;cellular systems;image analysis

298705

Basic Operations on the OTIS-Mesh Optoelectronic Computer.

AbstractIn this paper, we develop algorithms for some basic operationsbroadcast, window broadcast, prefix sum, data sum, rank, shift, data accumulation, consecutive sum, adjacent sum, concentrate, distribute, generalize, sorting, random access read and writeon the OTIS-Mesh [1] model. These operations are useful in the development of efficient algorithms for numerous applications [2].

Introduction The Optical Transpose Interconnection System ( OTIS ), proposed by Marsden et al. [4], is a hybrid optical and electronic interconnection system for large parallel computers. The OTIS architecture space optics to connect distant processors and electronic interconnect to connect nearby processors. Specifically, to maximize bandwidth, power efficiency, and to minimize system area and volume [1], the processors of an N 2 processor OTIS computer are partitioned into N groups of N processors each. Each processor is indexed by a tuple (G; G; G is the group index ( i.e., the group the processor is in ), and P the processor index within a group. The inter group interconnects are optical while the intra group interconnects are electronic. The optical or OTIS interconnects connect pairs of processors of the form [(G; P ); (P; G)]; that is, the group and processor indices are transposed by an optical interconnect. The electrical or intra group interconnections are according to any of the well studied electronic interconnection networks - mesh, hypercube, mesh of trees, and so forth. The choice if the electronic interconnection network defines a sub-family of OTIS computers - OTIS-Mesh, OTIS-Hypercube, and so forth. Figure 1 shows a 16 processor OTIS-Mesh. Each small square represents a processor. The number inside a processor square is the processor index P . Some processor squares have a pair (P The pair gives the row and column index of the processor P within its N \Theta N mesh. Each large This work was supported, in part, by the Army Research Office under grant DAA H04-95-1-0111. group 3 Figure 1: square encloses a group of processors. A group index G may also be given as a pair (G x ; G y ) where G x and G y are the row and column indices of the group assuming a N \Theta N layout of groups. Zane et al. [11] have shown that an N 2 processor OTIS-Mesh can simulate each move of a N \Theta N \Theta N \Theta four-dimensional ( 4D ) mesh computer using either one electronic move or one electronic and two OTIS moves ( depending on which dimension of the 4D mesh we are to move along ). They have also shown that an N 2 processor OTIS-Hypercube can simulate each move of an N 2 processor hypercube using either one electronic move or one electronic and two OTIS moves. Sahni and Wang [10, 9] have developed efficient algorithms to rearrange data according to bit- permute-complement permutations on OTIS-Mesh and OTIS-Hypercube computers, respectively. Rajasekaran and Sahni [7] have developed efficient randomized algorithms for routing, selection, and sorting on an OTIS-Mesh. In this paper, we develop deterministic OTIS-Mesh algorithms for the basic data operations for parallel computation that are studied in [8]. As shown in [8], algorithms for these operations can be used to arrive at efficient parallel algorithms for numerous applications, from image processing, computational geometry, matrix algebra, graph theory, and so forth. We consider both the synchronous SIMD and synchronous MIMD models. In both, all processors operate in lock-step fashion. In the SIMD model, all active processors perform the same operation in any step and all active processors move data along the same dimension or along OTIS connections. In the MIMD model, processors can perform different operations in the same step and can move data along different dimensions. 2 Basic Operations 2.1 Data Broadcast Data broadcast is, perhaps, the most fundamental operation for a parallel computer. In this operation, data that is initially in a single processor (G; P ) is to be broadcast or transmitted to all processors of the OTIS-Mesh. Data broadcast can be accomplished using the following three step algorithm: its data to all other processors in group G. Step 2: Perform an OTIS move. Step 3: Processor G of each group broadcasts the data within its group. Following Step 2, one processor of each group has a copy of the data, and following Step 3 each processor of the OTIS-Mesh has a copy. In the SIMD model, Steps 1 and 3 take 2( electronic moves each, and Step 2 takes one OTIS move. The SIMD complexity is 4( electronic moves and 1 OTIS move, or a total of 4 moves. Note that our algorithm is optimal because the diameter of the OTIS-Mesh is 4 example, if the data to be broadcast is initially in processor (0,0), the data needs to reach processor which is at a distance of 4 3. In the MIMD model, the complexity of Steps 1 and 3 depends on the value of ranges from a low of approximately to a high of 2( The overall complexity is at most 4( moves and one OTIS move. By contrast, simulating the 4D-mesh broadcast algorithm using the simulation method of [11] takes 4( electronic moves and 4( moves in the SIMD model and up to this many moves in the MIMD model. 2.2 Window Broadcast In a window broadcast, we start with data in the top left w \Theta w submesh of a single group G. Here w divides N . Following the window broadcast operation, the initial w \Theta w window tiles all groups; that is, the window is broadcast both within and across groups. Our algorithm for window broadcast is: Step 1: Do a window broadcast within group G. Step 2: Perform an OTIS move. Step 3: Do an intra group data broadcast from processor G of each group. Step 4: Perform an OTIS move. Following Step 1 the initial window properly tiles group G and we are left with the task of broadcasting from group G to all other groups. In Step 2, data d(G; P ) from (G; P ) is moved to In Step 3, d(G; P ) is broadcast to all processors moved to (i; P Step 1 of our window broadcast algorithm takes 2( moves in both the SIMD and MIMD models, and Step 3 takes 2( moves in the SIMD model and up to moves in the MIMD model. The total cost is 4 moves in the SIMD model and up to this many moves in the MIMD model. A simulation of the 4D mesh window broadcast algorithm takes the same number of electronic moves, but also takes 4( moves. 2.3 Prefix Sum The index (G; P ) of a processor may be transformed into a scalar I = GN+P with 0 - I ! N 2 . Let D(I) be the data in processor I, 0 - I ! N 2 . In a prefix sum, each processor I computes I . A simple prefix sum algorithm results from the following observation: where SD(I) is the sum of D(i) over all processors i that are in a group smaller than the group of I and LP (I) is the local prefix sum within the group of I. The simple prefix sum algorithm is: Step 1: Perform a local prefix sum in each group. Step 2: Perform an OTIS move of the prefix sums computed in Step 1 for all processors (G; N \Gamma 1). Step 3: Group modified prefix sum of the values, A, received in Step 2. In this modification, processor P computes rather than Step 4: Perform an OTIS move of the modified prefix sums computed in Step 3. Step 5: Each group does a local broadcast of the modified prefix sum received by its processor. Step Each processor adds the local prefix sum computed in Step 1 and the modified prefix sum it received in Step 5. The local prefix sums of Steps 1 and 3 take 3( moves in both the SIMD and MIMD models, and the local data broadcast of Step 5 takes 2( moves. The overall complexity is 8( moves and 2 OTIS moves. This can be reduced to moves and 2 OTIS moves by deferring some of the Step 1 moves to Step 5 as below. Step 1: In each group, compute the row prefix sums R. Step 2: Column of each group computes the modified prefix sums of its R values. Step 3: Perform an OTIS move on the prefix sums computed in Step 2 for all processors (G; N \Gamma 1). Step 4: Group modified prefix sum of the values, A, received in Step 3. Step 5: Perform an OTIS move of the modified prefix sums computed in Step 4. Step Each group broadcasts the modified prefix sum received in Step 5 along column of its mesh. Step 7: The column processors add the modified prefix sum received in Step 6 and the prefix sum of R values computed in Step 2 minus its own R value computed in Step 1. Step 8: The result computed by column processors in Step 7 is broadcast along mesh rows. Step 9: Each processor adds its R value and the value it received in Step 8. If we simulate the best 4D mesh prefix sum algorithm, the resulting OTIS mesh algorithm takes moves. 2.4 Data Sum In this operation, each processor is to compute the sum of the D values of all processors. An optimal SIMD data sum algorithm is: Step 1: Each group performs the data sum. Step 2: Perform an OTIS move. Step 3: Each group performs the data sum. In the SIMD model Steps 1 and 3 take 4( moves, and step 2 takes 1 OTIS move. The total cost is 8( moves. Note that since the distance between processors (0; moves and since each needs to get information from the other, at least 8( are needed ( the moves needed to send information from (0; 0) to (N \Gamma and those from cannot be overlapped in the SIMD model ). Also, note that a simulation of the 4D mesh data sum algorithm takes 8( moves. The MIMD complexity can be reduced by computing the group sums in the middle processor of each group rather than in the bottom right processor. The complexity now becomes 4( electronic and 1 OTIS moves when N is odd and 4 N electronic and 1 OTIS moves when is even. The simulation of the 4D mesh, however, takes 4( moves. Notice that the MIMD algorithm is near optimal as the diameter of the OTIS-Mesh isp 2.5 Rank In the rank operation, each processor I has a flag S(I) 2 f0; 1g, 0 - I ! N 2 . We are to compute the prefix sums of the processors with This operation can be performed in 7( electronic and 2 OTIS moves using the prefix sum algorithm of Section 2.3. 2.6 Shift Although there are many variations of the shift operation, the ones we believe are most useful in application development are: (a) mesh row shift with zero fill - in this we shift data from processor (G x N . The shift is done with zero fill and end discard ( i.e., if or P y the data from P y is discarded ). (b) mesh column shift with zero fill - similar to (a), but along mesh column P x . (c) circular shift on a mesh row - in this we shift data from processor (G x (d) circular shift on a mesh column - similar to (c), but instead P x is used. row shift with zero fill - similar to (a), except that G y is used in place of P y . (f) group column shift with zero fill - similar to (e), but along group column G x . circular shift on a group row - similar to (c), but with G y rather than P y . circular shift on a group column - similar to (g), with G x in place of G y . Shifts of types (a) through (d) are done using the best mesh algorithms while those of types (e) through (h) are done as below: 1: Perform an OTIS move. Step 2: Do the shift as a P x ( if originally a G x shift ) or a P y ( if originally a G y shift ) shift. Step 3: Perform an OTIS move. Shifts of types (a) and (b) take s electronic moves on the SIMD and MIMD models; (c) and (d) take electronic moves on the SIMD model and maxfjsj; moves on the MIMD model; (e) and (f) take s electronic and 2 OTIS moves on both SIMD and MIMD models; and (g) and (h) take N electronic and 2 OTIS moves on the SIMD model and maxfjsj; electronic and 2 OTIS moves on the MIMD model. If we simulate the corresponding 4D mesh algorithms, we obtain the same complexity for (a) - (d), but (e) and (f) take an additional 2s \Gamma 2 OTIS moves, and (g) and (h) take an additional 2 \Theta maxfjsj; moves. 2.7 Data Accumulation Each processor is to accumulate M , , values from its neighboring processors along one of the four dimensions G x , G y , be the data in processor In a data accumulation along the G x dimension ( for example ), each processor accumulates in an array A the data values from ((G x Specifically, we have Accumulation in other dimensions is similar. The accumulation operation can be done using a circular shift of \GammaM in the appropriate dimen- sion. The complexity is readily obtained from that for the circular shift operation ( see Section 2.6 2.8 Consecutive Sum The N 2 processor OTIS-Mesh is tiled with one-dimensional blocks of size M . These blocks may align with any of the four dimensions G x , G y , P x , and P y . Each processor has M values X[j], . The ith processor in a block is to compute the sum of the X[i]s in that block. Specifically, processor i of a block computes where i and j are indices relative to a block. When the one-dimensional blocks of size M align with the P x or P y dimensions, a consecutive sum can be performed by using M tokens in each block to accumulate the M sums S(i), Assume the blocks align along P x . Each processor in a block initiates a token labeled with the processor's intra block index. The tokens from processors 0 through are right bound and that from M \Gamma 1 is left bound. In odd time steps, right bound tokens move one processor right along the block, and in even time steps left bound tokens move one processor left along the block. When a token reaches the rightmost or leftmost processor in the block, it reverses direction. Each token visits each processor in its block twice - once while moving left and once while moving right. During the rightward visits it adds in the appropriate X value from the processor. After time steps ( and hence moves ), all tokens return to their originating processors, and we are done. In the MIMD model, the left and right moves can be done simultaneously, and only electronic moves are needed. When the one-dimensional size M blocks align with G x or G y , we first do an OTIS move; then run either a P x or P y consecutive sum algorithm; and then do an OTIS move. The number of electronic moves is the same as for P x or P y alignment. However, two additional OTIS moves are needed. Simulation of the corresponding 4D mesh algorithm takes an additional for the case of G x or G y alignment in the SIMD model and an additional moves in the MIMD model. 2.9 Adjacent Sum This operation is similar to the data accumulation operation of Section 2.7 except that the M accumulated values are to be summed. The operation can be done with the same complexity as data accumulation using a similar algorithm. 2.10 Concentrate A subset of the processors contain data. These processors have been ranked as in Section 2.5. So the data is really a pair (D; r); D is the data in the processor and r is its rank. Each pair (D; r) is to be moved to processor r, 0 - r ! b, where b is the number of processors with data. Using the (G; P ) format for a processor index, we see that (D; r) is to be routed from its originating processor to processor (br=Nc; r mod N ). We accomplish this using the steps: Step 1: Each pair (D; r) is routed to processor r mod N within its current group. Step 2: Perform an OTIS move. Step 3: Each pair (D; r) is routed to processor br=Nc within its current group. Step 4: Perform an OTIS move. Theorem 1 The four step algorithm given above correctly routes every pair (D; r) to processor Proof Step 1 does the routing on the second coordinate. This step does not route two pairs to the same processor provided no group has two pairs (D Since each group has at most N pairs and the ranks of these pairs are contiguous integers, no group can have two pairs with r 1 mod each processor has at most one pair and each pair is in the correct processor of the group, though possibly in the wrong group. To get the pairs to their correct groups without changing the within group index, Step 2 performs an OTIS move, which moves data from processor (G; P ) to processor (P; G). Now all pairs in a group have the same r mod N value and different br=Nc values. The routing on the br=Nc values, as in Step 3, routes at most one pair to each processor. The OTIS move of Step 4, therefore, gets every pair to its correct destination processor. 2 In group 0, Step 1 is a concentrate localized to the group, and in the remaining groups, Step 1 is a generalized concentrate in which the ranks have been increased by the same amount. In all groups we may use the mesh concentrate algorithm of [6] to accomplish the routing in 4( moves. Step 3 is also a concentrate as the br=Nc values of the pairs are in ascending order from 0; moves each in the SIMD model and in the MIMD model [6]. Therefore, the overall complexity of concentrate is 8( electronic and 2 OTIS moves in the SIMD model and 4( moves in the MIMD model. We can improve the SIMD time to 7( moves by using a better mesh concentrate algorithm than the one in [6]. The new and simpler algorithm is given below for the case of a generalized concentration on a N \Theta mesh. Step 1: Move data that is to be in a column right of the current one rightwards to the proper processor in the same row. Step 2: Move data that is to be in a column left of the current one leftwards to the proper processor in the same row. Step 3: Move data that is to be in a smaller row upwards to the proper processor in the same column. Step 4: Move data that is to be in a bigger row downwards to the proper processor in the same column. In a concentrate operation on a square mesh data that begins in two processors of the same row ends up in different columns as the rank of these two data differs by at most and 2 do not leave two or more data in the same processor. Steps 3 and 4 get data to the proper row and hence to the proper processor. Note that it is possible to have up to two data items in a processor following Step 1 and Step 3. The complexity of the above concentrate algorithm is on a SIMD mesh and 2( on an MIMD mesh ( we can overlap Steps 1 and 2 as well as Steps 3 and 4 on an MIMD mesh ). For an ordinary concentrate in which the ranks begin at 1, Step 4 can be omitted as no data moves down a column to a row with bigger index. So an ordinary concentrate takes only 3( moves. This improves the SIMD concentration algorithm of [6], which takes 4( moves to do an ordinary concentrate. Actually, we can show that the four step concentration algorithm just stated is optimal for the SIMD model. Consider the ordinary concentrate instance in which the selected elements are in processors (0; 0). The ranks are 0, 1, \Delta \Delta \Delta, 1. So the data in processor (0; is to be moved to processor (0,0). This requires moves that yield a net of moves. Also, the data in processor ( is to be moved to processor (0; This requires a net of moves and moves. None of these moves can be overlapped in the SIMD model. So every SIMD concentrate algorithm must take at least moves in each of the directions left, right, and up; a total of at least 3( moves. For the generalized concentrate algorithm, the ranks need not start at zero. Suppose we have two elements to concentrate. One is at processor (0,0) and has rank N \Gamma 1, and the other is at processor ( has rank N . The data in (0,0) is to be moved to ( at a cost of right and down moves. The data in ( is to be moved to (0,0) at a cost of net left and up moves. So at least 4( are needed. Theorem 2 The OTIS-Mesh data concentration algorithm described above is optimal for both the SIMD and MIMD models; that is, (a) every SIMD concentration algorithm must make 7( electronic and 2 OTIS moves in the worst case, and (b) every MIMD concentration algorithm must make 4( moves. Proof (a) Suppose that the data to be concentrated are in the processors shown in Table 1. Let a denote processor ( and let c denote processor (0,1,0,0). The ranks of a, b, and c are N 3=2 , N 3=2 respectively. Therefore, following the concentration the data D(a), D(b), and D(c) initially in processors a, b, and c will be in processors (0,1,0,0), (0; respectively. Figure 2 shows the initial and concentrated data layout for the case when The change in G x , G y , P x , and P y values between the final and initial locations of D(a), D(b), and D(c) is shown in Table 2. a c x x x x x x x x x x x x x x x x x x x x x x x x x x x x (b) a c x x x x x x x x x x x x x x x x x x x x x x x x x x (a) Figure 2: Data Configuration: (a) Initial; (b) Concentrated Table 1: Processors with data to concentrate data G x G y D(a) \Gamma( Table 2: Net change in G x , G y , P x , and P y The maximum net negative change in each of G x , G y , P x , and P y is \Gamma( 1). Since a net negative change in G x can only be overlapped with a net negative change in P x and since D(b) needs \Gamma( negative change in both G x and P x , we must make at least 2( moves that decrease the row index within a mesh. Similarly, because of D(a)'s requirements, at least 2( moves that increase the column index within a N \Theta mesh must be made. Turning our attention to net positive changes, we see that because of D(b)'s requirements there must be at least 2( moves that increase the column index. D(c) requires moves that increase the row index. Since positive net moves cannot be overlapped with negative net moves, and since net moves along G x and P x cannot be overlapped with net moves along G y and P y , the concentration of the configuration of Table 1 must take at least 7( moves. In addition to 7( moves, we need at least 2 OTIS moves to concentrate the data of Table 1. To see this consider the data initially in group (0,1). This data is in group (0,0) following the concentration. At least one OTIS move is needed to move the data out of group (0,1). A nontrivial OTIS-Mesh has - 2 processors on a row of a N \Theta N submesh. For such an OTIS-Mesh, at least two pieces of data must move from group (0,1) to group (0,0). A single OTIS move scatters data from group (0,1) to different groups with each data going to a different group. At least one additional OTIS move must be made to get the data back into the same group. Therefore the concentration of the configuration of Table 1 cannot be done with fewer than 2 OTIS moves. (b) Consider the initial configuration of Table 1. Since the shortest path between processor b and its destination processor is 4( and one OTIS move, at least that many electronic moves are made, in the worst case, by every concentration algorithm. The reason that at least 2 OTIS moves are needed to complete the concentration is the same as for (a). 2 2.11 Distribute This is the inverse of the concentrate operation of Section 2.10. We start with pairs (D in the first q +1 processors 0; and are to route pair (D i ; d i ) to processor q. The algorithm of Section 2.10 tells us how to start with pairs (D i ; i) in processor move them so that D i is in i. By running this backwards, we can start with D i in i and route it to d i . The complexity of the distribute operation is the same as that of the concentrate operation. We have shown that the concentrate algorithm of Section 2.10 is optimal; it follows that the distribute algorithm is also optimal. 2.12 Generalize We start with the same initial configuration as for the distribute operation. The objective is to have D i in all processors j such that d i we simulate the 4D mesh algorithm for generalize using the simulation strategy of [11], it takes 8( and 8( moves to perform the generalize operation on an SIMD OTIS-Mesh. We can improve this to 8( moves if we run the generalize algorithm of [6] adapted to use OTIS moves as necessary. The outer loop of the algorithm of [6] examines processor index bits from 2p \Gamma 1 to 0 where . So in the first p iterations we are moving along bits of the G index and in the last p iterations along bits of the P index. On an OTIS-Mesh we would break this into two parts as below: 1: Perform an OTIS move. Step 2: Run the GENERALIZE procedure of [6] from bit while maintaining the original index. Step 3: Perform an OTIS move. Step 4: Run the GENERALIZE algorithm of [6] from bit On an MIMD OTIS-Mesh the above algorithm takes 4( moves. We can reduce the SIMD complexity to 7( moves by using a better algorithm to do the generalize operation on a 2D SIMD mesh. This algorithm uses the same observation as used by us in Section 2.10 to speed the 2D SIMD mesh concentrate algorithm; that is, of the four possible move directions, only three are possible. When doing a generalize on a 2D N \Theta N mesh the possible move directions for data are to increasing row indexes and to decreasing and increasing column indexes. With this observation, the algorithm to generalize on a 2D mesh becomes: Step 1: Move data along columns to increasing row indexes if the data is needed in a row with higher index. Step 2: Move data along rows to increasing column indexes if the data is needed in a processor in that row with higher column index. Step 3: Move data along rows to decreasing column indexes if the data is needed in a processor in that row with smaller column index. The correctness of the preceding generalize algorithm can be established using the argument of Theorem 1, and its optimality follows from Theorem 2 and the fact that the distribute operation, which is the inverse of the concentrate operation, is a special case of the generalize operation. The new and more efficient generalize algorithm may be used in Step 2 of the OTIS-Mesh generalize algorithm. It cannot be used in Step 4 because the generalize of this step requires the full capability of the code of [6] which permits data movement in all four directions of a mesh. When we use the new generalize algorithm for Step 2 of the OTIS-Mesh generalize algorithm, we can perform a generalize on a SIMD OTIS-Mesh using 7( moves. The new algorithm is optimal for both SIMD and MIMD models. This follows from the lower bound on a concentrate operation established in Theorem 2 and the observation made above that the distribute operation, which is a special case of the generalize operation, is the inverse of the concentrate operation and so has the same lower bound. 2.13 Sorting As was the case for the operations considered so far, an O( time algorithm to sort can be obtained by simulating a similar complexity 4D mesh algorithm. For sorting a 4D Mesh, the Figure 3: Row-Column Transformation of Leighton's Column Sort algorithm of Kunde [2] is the fastest. Its simulation will sort into snake-like row-major order usingp N) electronic and 12 OTIS moves on the SIMD model and 7 electronic and 6 OTIS moves on the MIMD model. To sort into row-major order, additional moves to reverse alternate dimensions are needed. This means that an OTIS-Mesh simulation of Kunde's 4D mesh algorithm to sort into row-major order will take electronic and 16 OTIS moves on the SIMD model. We show that Leighton's column sort [3] can be implemented on an OTIS-Mesh to sort into row-major order using 22 electronic and O(N 3=8 ) OTIS moves on the SIMD model and 11 N) electronic and O(N 3=8 ) OTIS moves on the MIMD model. Our OTIS-Mesh sorting algorithm is based on Leighton's column sort [3]. This sorting algorithm sorts an r \Theta s array, with r - using the following seven steps: Step 1: Sort each column. Step 2: Perform a row-column transformation. Step 3: Sort each column. Step 4: Perform the inverse transformation of Step 2. Step 5: Sort each column in alternating order. Step Apply two steps of comparison-exchange to adjacent rows. Step 7: Sort each column. Figure 3 shows an example of the transformation of Step 2, and its inverse. Figure 4 shows a step by step example of Leighton's column sort. \Gamma! 11 \Gamma! Figure 4: Example of Leighton's Column Sort Although Leighton's column sort is explicitly stated for r \Theta s arrays with r - can be used to sort arrays with s - into row-major order by interchanging the roles of rows and columns. We shall do this and use Leighton's method to sort an N 1=2 \Theta N 3=2 array. We interpret our N 2 OTIS-Mesh as an N 1=2 \Theta N 3=2 array with G x giving the row index and G y giving the column index of an element processor. We shall further subdivide G x ( G y , P x , P y , and G x 4 from left to right. We use G x 2\Gamma4 , for example, to . Since bits and G x i has p=8 bits. These notations are helpful in describing the transformations in Steps 2 and 4 of the column sort, as we use the BPC permutations of [5] to realize these transformations. A BPC permutation [5] is specified by a vector (a) A i 2 f\Sigma0; (b) [jA is a permutation of [0; The destination for the data in any processor may be computed in the following manner. Let be the binary representation of the processor's index. Let d be that of the destination processor's index. Then, In this definition, \Gamma0 is to be regarded as ! 0, while +0 is - 0. Table 3 shows an example of the BPC permutation defined by the permutation vector \Gamma3] on a 16 processor OTIS-Mesh. Source Destination Processor (G; P ) Binary Binary (G; P ) Processor 9 (2,1) 1001 0000 (0,0) 0 Table 3: Source and destination of the BPC permutation [\Gamma0; 1; 2; \Gamma3] in a 16 processor OTIS- Mesh In describing our sorting algorithm, we shall, at times, use a 4D array interpretation of an OTIS-Mesh. In this interpretation, processor of the OTIS-Mesh corresponds to processor of the 4D mesh. We use g x to denote the bit positions of G x , that is the leftmost p=2 bits in a processor index, g x1 to represent the leftmost p=8 bit positions, p y to represent the rightmost p=2 bit positions, p y 3\Gamma4 to represent the rightmost p=4 bit positions, and so on. Our strategy for the sorting steps 1, 3, 5, and 7 of Leighton's method is to collect each row ( recall that since we are sorting an N 1=2 \Theta N 3=2 array, the column-sort steps of Leighton's method become row-sort steps ) of our N 1=2 \Theta N 3=2 array into an N 3=8 \Theta N 3=8 \Theta N 3=8 \Theta N 3=8 4D submesh of the OTIS-Mesh, and then sort this row by simulating the 4D mesh sort algorithm of [2]. This strategy translates into the following sorting algorithm: rows of the N 1=2 \Theta N 3=2 array into N 3=8 \Theta N 3=8 \Theta N 3=8 \Theta N 3=8 4D submeshes Perform the BPC permutation P 2: [ Sort each row of the N 1=2 \Theta N 3=2 array Sort each 4D submesh of size N 3=8 \Theta N 3=8 \Theta N 3=8 \Theta N 3=8 . 3: [ Do the inverse of Step 1, perform a column-row transformation, and move rows into 3=8 \Theta N 3=8 \Theta N 3=8 \Theta N 3=8 submeshes Perform the BPC permutation P each row of the N 1=2 \Theta N 3=2 array Sort each 4D submesh of size N 3=8 \Theta N 3=8 \Theta N 3=8 \Theta N 3=8 . 5: [ Do the inverse of Step 1, perform a row-column transformation, and move rows into 3=8 \Theta N 3=8 \Theta N 3=8 \Theta N 3=8 submeshes Perform the BPC permutation P 0 x 1\Gamma3 each row in alternating order ] Sort each 4D submesh of size N 3=8 \Theta N 3=8 \Theta N 3=8 \Theta N 3=8 . 7: [ Move rows back from 4D submeshes Perform the BPC permutation P 0 Step 8: Apply two steps of comparison-exchange to adjacent rows. into submeshes of size N 3=8 \Theta N 3=8 \Theta N 3=8 \Theta N 3=8 ] Perform the BPC permutation P each row of the N 1=2 \Theta N 3=2 array Sort each 4D submesh of size N 3=8 \Theta N 3=8 \Theta N 3=8 \Theta N 3=8 . rows back from 4D submeshes Perform the BPC permutation P 0 y 2\Gamma4 Notice that the row to 4D submesh transform is accomplished by the BPC permutation P y 2\Gamma4 Elements in the same row of our N 1=2 \Theta N 3=2 array interpretation have the same G x value; but in our 4D mesh interpretation, elements in the same 3=8 \Theta N 3=8 \Theta N 3=8 \Theta N 3=8 submesh have the same G x 1 value. P a results in this prop- erty. To go from Step 2 to Step 3 of Leighton's method, we need to first restore the N 1=2 \Theta N 3=2 array interpretation using the inverse permutation of P a , that is, perform the BPC permutation y 2\Gamma4 ]; then perform a column-row transform using BPC permutation finally map the rows of our N 1=2 \Theta N 3=2 array into 4D submeshes of size N 3=8 \ThetaN 3=8 \ThetaN 3=8 \ThetaN 3=8 using the BPC permutation P a . The three BPC permutation sequence a a is equivalent to the single BPC permutation P The preceding OTIS-Mesh implementation of column sort performs 6 BPC permutations, 4 4D mesh sorts, and two steps of comparison-exchange on adjacent rows. Since the sorting steps take O(N 3=8 ) time each ( use Kunde's 4D mesh sort [2] followed by a transform from snake-like row-major to row-major ), and since the remaining steps take O(N 1=2 ) time, we shall ignore the complexity of the sort steps. We can reduce the number of BPC permutations from 6 to 3 as follows. First note that the P a of Step 1 just moves elements from rows of the N 1=2 \Theta N 3=2 array into N 3=8 \Theta N 3=8 \Theta N 3=8 \Theta N 3=8 4D submeshes. For the sort of Step 2, it doesn't really matter which N 3=2 elements go to each 4D submesh as the initial configuration is an arbitrary unsorted configuration. So we may eliminate note that the BPC permutations of Steps 7 and 9 cancel each other and we can perform the comparison-exchange of Step 8 by moving data from one N 3=8 \Theta N 3=8 \Theta N 3=8 \Theta N 3=8 4D submesh to an adjacent one and back in O(N 3=8 ) time. With these observations, the algorithm to sort on an OTIS-Mesh becomes: 1: Sort in each subarray of size N 3=8 \Theta N 3=8 \Theta N 3=8 \Theta N 3=8 Step 2: Perform the BPC permutation P c . Step 3: Sort in each subarray. Step 4: Perform the BPC permutation P 0 c . Step 5: Sort in each subarray. Step Apply two steps of comparison-exchange to adjacent subarrays. Step 7: Sort in each subarray. Step 8: Perform the BPC permutation P 0 a . Using the BPC routing algorithm of [10], the three BPC permutations can be done usingp N electronic and 3 log moves on the SIMD model and N electronic and moves on the MIMD model. A more careful analysis based on the development in [5] and [10] reveals that the permutations P 0 a , P c , and P 0 c can be done with 28 N electronic and log moves on the SIMD model and 14 N electronic and 3 log on the MIMD model. By using p 0 y 2\Gamma4 ], the permutation cost becomes 22 N electronic and log 2 N+5 OTIS moves on the SIMD model and 11 and log 2 N+5 OTIS moves on the MIMD model. The total number of moves is thus 22 electronic and O(N 3=8 ) OTIS moves on the SIMD model and 11 O(N 3=8 ) OTIS moves on the MIMD model. This is superior to the cost of the sorting algorithm that results from simulating the 4D row-major mesh sort of Kunde [2]. 2.14 Random Access Read ( RAR ) In a random access read (RAR) [8] processor I wishes to read data variable D of processor d I , . The steps suggested in [8] for this operation are: Step 0: Processor I creates a triple (I; D; d I ) where D is initially empty. 1: Sort the triples by d I . Step 2: Processor I checks processor I +1 and deactivates if both have triples with the same third coordinate. Step 3: Rank the remaining processors. Step 4: Concentrate the triples using the ranks of Step 3. Step 5: Distribute the triples according to their third coordinates. Step Load each triple with the D value of the processor it is in. Step 7: Concentrate the triples using the ranks in Step 3. Step 8: Generalize the triples to get the configuration we had following Step 1. Step 9: Sort the triples by their first coordinates. Using the SIMD model the RAR algorithm of [8] take 79( moves and O(N 3=8 ) OTIS moves. On the MIMD model, it takes 45( moves. 2.15 Random Access Write ( RAW ) Now processor I wants to write its D data to processor d I , 0 - I ! N 2 . The steps in the RAW algorithm of [8] are: Step 0: Processor I creates the tuple (D(I); d I 1: Sort the tuples by their second coordinates. Step 2: Processor I deactivates if the second coordinate of its tuple is the same as the second coordinate of the tuple in I Step 3: Rank the remaining processors. Step 4: Concentrate the tuples using the ranks of Step 3. Step 5: Distribute the tuples according to their second coordinates. implements the arbitrary write method for a concurrent write. In this, any one of the processors wishing to write to the same location is permitted to succeed. The priority model may be implemented by sorting in Step 1 by d I and within d I by priority. The common and combined models can also be implemented, but with increased complexity. On the SIMD model, an RAW takes 43( moves while on the MIMD model, it takes 26( moves. 3 Conclusion We have developed OTIS-Mesh algorithms for the basic parallel computing algorithms of [8]. Our algorithms run faster than the simulation of the fastest algorithms known for 4D meshes. Table 4 summarizes the complexities of our algorithms and those of the corresponding ones obtained by simulating the best 4D-mesh algorithms. Note that the worst case complexities are listed for the broadcast and window broadcast operation, and that of the case when N is even is presented for the data sum operation on the MIMD model. Also, the complexities listed for circular shift, data accumulation, and adjacent sum assume that the shift distance is - N=2 on the MIMD model. Table 4 gives only the dominating N terms for sorting. Our algorithms for data broadcast, data sum, concentrate, distribute, and generalize are optimal. --R Routing and sorting on mesh-connected arrays Tight bounds on the complexity of parallel sorting. Optical transpose interconnection system architectures. An optimal routing algorithm for mesh-connected parallel computers Data broadcasting in SIMD computers. Randomized routing Hypercube Algorithms with Applications to Image processing and Pattern Recognition. BPC permutations on the OTIS-Hypercube optoelectronic computer BPC permutations on the OTIS-Mesh optoelectronic computer Scalable network architectures using the optical transpose interconnection system (OTIS). --TR --CTR A. Al-Ayyoub , A. Awwad , K. Day , M. Ould-Khaoua, Generalized methods for algorithm development on optical systems, The Journal of Supercomputing, v.38 n.2, p.111-125, November 2006 Behrooz Parhami, The Hamiltonicity of swapped (OTIS) networks built of Hamiltonian component networks, Information Processing Letters, v.95 n.4, p.441-445, 31 August 2005 Ahmad M. Awwad, OTIS-star an attractive alternative network, Proceedings of the 4th WSEAS International Conference on Software Engineering, Parallel & Distributed Systems, p.1-6, February 13-15, 2005, Salzburg, Austria Khaled Day , Abdel-Elah Al-Ayyoub, Topological Properties of OTIS-Networks, IEEE Transactions on Parallel and Distributed Systems, v.13 n.4, p.359-366, April 2002 Xiaofan Yang , Graham M. Megson , David J. Evans, An oblivious shortest-path routing algorithm for fully connected cubic networks, Journal of Parallel and Distributed Computing, v.66 n.10, p.1294-1303, October 2006 Behrooz Parhami, Swapped interconnection networks: topological, performance, and robustness attributes, Journal of Parallel and Distributed Computing, v.65 n.11, p.1443-1452, November 2005 Ahmad M. Awwad, OTIS-star an attractive alternative network, Proceedings of the 4th WSEAS International Conference on Software Engineering, Parallel & Distributed Systems, p.1-6, February 13-15, 2005, Salzburg, Austria Khaled Day, Optical transpose k-ary n-cube networks, Journal of Systems Architecture: the EUROMICRO Journal, v.50 n.11, p.697-705, November 2004 Prasanta K. Jana, Polynomial interpolation and polynomial root finding on OTIS-mesh, Parallel Computing, v.32 n.4, p.301-312, April 2006 Chih-fang Wang , Sartaj Sahni, Matrix Multiplication on the OTIS-Mesh Optoelectronic Computer, IEEE Transactions on Computers, v.50 n.7, p.635-646, July 2001 Chih-Fang Wang , Sartaj Sahni, Image Processing on the OTIS-Mesh Optoelectronic Computer, IEEE Transactions on Parallel and Distributed Systems, v.11 n.2, p.97-109, February 2000

optoelectronic;random access write;distribute;adjacent sum;random access read;prefix sum;OTIS-Mesh;concentrate;data accumulation;window broadcast;consecutive sum;sorting;broadcast;data sum;generalize;shift

298770

Co-Evolution in the Successful Learning of Backgammon Strategy.

Following Tesauros work on TD-Gammon, we used a 4,000 parameter feedforward neural network to develop a competitive backgammon evaluation function. Play proceeds by a roll of the dice, application of the network to all legal moves, and selection of the position with the highest evaluation. However, no backpropagation, reinforcement or temporal difference learning methods were employed. Instead we apply simple hillclimbing in a relative fitness environment. We start with an initial champion of all zero weights and proceed simply by playing the current champion network against a slightly mutated challenger and changing weights if the challenger wins. Surprisingly, this worked rather well. We investigate how the peculiar dynamics of this domain enabled a previously discarded weak method to succeed, by preventing suboptimal equilibria in a meta-game of self-learning.

Introduction It took great chutzpah for Gerald Tesauro to start wasting computer cycles on temporal difference learning in the game of Backgammon (Tesauro, 1992). Letting a machine learning program play itself in the hopes of becoming an expert, indeed! After all, the dream of computers mastering a domain by self-play or "introspection" had been around since the early days of AI, forming part of Samuel's checker player (Samuel, 1959) and used in Donald Michie's MENACE tic-tac-toe learner (Michie, 1961); but such self-conditioning systems had later been generally abandoned by the field due to problems of scale and weak or non-existent internal representations. Moreover, self- playing learners usually develop eccentric and brittle strategies which appear clever but fare poorly against expert human and computer players. Yet Tesauro's 1992 result showed that this self-play approach could be powerful, and after some refinement and millions of iterations of self-play, his TD-Gammon program has become one of the best backgammon players in the world (Tesauro, 1995). His derived weights are viewed by his corporation as significant enough intellectual property to keep as a trade secret, except to leverage sales of their minority operating system (International Business Machines, 1995). Others have replicated this TD result in backgammon both for research purposes (Boyan, 1992) and commercial purposes. While reinforcement learning has had limited success in other areas (Zhang and Dietterich, 1996, Crites and Barto, 1996, Walker et al., 1994), with respect to the goal of a self-organizing learning machine which starts from a minimal specification and rises to great sophistication, TD-Gammon stands alone. How is its success to be understood, explained, and replicated in other domains? Our hypothesis is that the success of TD-gammon is not principally due to the back-propagation, reinforcement, or temporal-difference technologies, but to an inherent bias from the dynamics of the game of backgammon, and the co-evolutionary setup of the training, by which the task dynamically changes as the learning progresses. We test this hypothesis by using a much simpler co-evolutionary learning method for backgammon namely hill-climbing. 2. Implementation Details We use a standard feedforward neural network with two layers and the sigmoid set up in the same fashion as (Tesauro, 1992) with 4 units to represent the number of each player's pieces on each of the 24 points, plus 2 units each to indicate how many are on the bar and off the board. In addition, we added one more unit which reports whether or not the game has reached the endgame or "race" situation, making a total of 197 input units. These are fully connected to 20 hidden units, which are then connected to one output unit that judges the position. Including bias on the hidden units, this makes a total of 3980 weights. The game is played by generating all legal moves, converting them into the proper network input, and picking the position judged as best by the network. We started with all weights set to zero. Our initial algorithm was hillclimbing: 1. add gaussian noise to the weights 2. play the network against the mutant for a number of games 3. if the mutant wins more than half the games, select it for the next generation. The noise was set so each step would have a 0.05 RMS distance (which is the euclidean distance divided by ). Surprisingly, this worked reasonably well. The networks so evolved improved rapidly at first, but then sank into mediocrity. The problem we perceived is that comparing two close backgammon players is like tossing a biased coin repeatedly: it may take dozens or even hundreds of games to find out for sure which of them is better. Replacing a well-tested champion is dangerous without enough information to prove the challenger is really a better player and not just a lucky novice. Rather than burden the system with so much computation, we instead introduced the following modifications to the algorithm to avoid this "Buster Douglas Effect": 2 Firstly, the games are played in pairs, with the order of play reversed and the same random seed used to generate the dice rolls for both games. This washes out some of the unfairness due to the dice rolls when the two networks are very close - in particular, if they were identical, the result would always be one win each - though, admittedly, if they make different moves early in the game, what is a good dice roll at a particular move of one game may turn out to be a bad roll at the corresponding move of the parallel game. Secondly, when the challenger wins the contest, rather than just replacing the champion by the challenger, we instead make only a small adjustment in that direction: champion This idea, similar to the "inertia" term in back-propagation (Rumelhart et al., 1986) was introduced on the assumption that small changes in weights would lead to small changes in decision-making by the evaluation function. So, by just "biting the ear" off the challenger and adding it to the champion, most of the current decisions are preserved , and we would be less likely to have a catastrophic replacement of the champion by a lucky novice challenger. In the initial stages of evolution, two pairs of parallel games were played and the challenger was required to win 3 out of 4 of these games. Although we would have liked to rank our players against the same players used - Neurogammon and Gammontool - these were not available to us. Figure 1 shows the first 35,000 players rated against PUBEVAL, a moderately good public-domain player trained by Tesauro using human expert preferences. There are three things to note: (1) the percentage of wins against PUBEVAL increases from 0% to about 33% by 20,000 generations, (2) the frequency of successful challengers increases over time as the player improves, and (3) there are epochs (e.g. starting at 20,000) where the performance against PUBEVAL begins to falter. The first fact shows that our simple self- 2. Buster Douglas was world heavyweight boxing champion for 9 months in 1990. playing hill-climber is capable of learning. The second fact is quite counter-intuitive - we expected that as the player improved, it would be harder to challenge it! This is true with respect to a uniform sampling of the 4000 dimensional weight space, but not true for a sampling in the neighborhood of a given player: once the player is in a good part of weight space, small changes in weights can lead to mostly similar strategies, ones which make mostly the same moves in the same situations. However, because of the few games we were using to determine relative fitness, this increased rate of change allows the system to drift, which may account for the subsequent degrading of performanceTo counteract the drift, we decided to change the rules of engagement as the evolution proceeds according to the following "annealing schedule": after 10,000 generations, the number of games that the challenger is required to win was increased from 3 out of 4 to 5 out of 6; after 70,000 generations, it was further increased to 7 out of 8 (of course each bout was abandoned as soon as the champion won more than one game, making the average number of games per generation considerably less than 8). The numbers 10,000 and 70,000 were chosen on an ad hoc basis from observing the frequency of successful challenges and the Buster Douglas effect in this particular run, but later experiments showed how to determine the annealing schedule in a more principled manner (see Section 3.2 below). After 100,000 games using this simple hill-climb, we have developed a surprising player, capable of winning 40% of the games against PUBEVAL. The networks were sampled every 100 generations in order to test their performance. Networks at generation 1,000, 10,000 and 100,000 were extracted and used as benchmarks. Figure 2 shows the percentage of wins for the sampled players against the three benchmark networks. Note that the three curves cross the 50% line at 1, 10, and 100, respectively and show a general improvement over time. The end-game of backgammon, called the "bear-off," can be used as another yardstick of the progress of learning. The bear-off occurs when all of a player's pieces are in their home board, or first 6 points, and then the dice rolls can be used to remove pieces Figure 1: Percentage of wins of our first 35,000 generation players against PUBEVAL. Each match consisted of 200 games. Generation %win from the board. To test our network's ability at the end-game, we set up a racing board with two pieces on each player's 1 through 7 point and one piece on the 8 point. The graph in Figure 3 shows the average number of rolls to bear-off for each network playing itself using a fixed set of 200 random dice-streams. We note that PUBEVAL is stronger at 16.6 rolls, and will discuss its strengths and those of Tesauro's 1992 results in Section 5. Figure 2: Percentage of wins against benchmark networks 1,000 [upper], 10,000 [middle] and 100,000 [lower]. This shows a noisy but nearly monotonic increase in player skill as evolution proceeds. %win Generation Generation Figure 3: Average number of rolls to bearoff by each generation, sampled with 200 dice-streams. PUBEVAL averaged 16.6 rolls for the task. 3. Analysis 3.1. Learnability and Unlearnability Learnability can be formally defined as a time constraint over a search space. How hard is it to randomly pick 4000 floating-point weights to make a good backgammon evaluator? It is simply impossible. How hard is it to find weights better than the current Initially, when all weights are random, it is quite easy. As the playing improves, we would expect it to get harder and harder, perhaps similar to the probability of a tornado constructing a 747 out of a junkyard. However, if we search in the neighborhood of the current weights, we will find many similar players which make mostly the same moves but which can capitalize on each other's slightly different choices and exposed weaknesses in a tournament. Note that this is a different point than Tesauro originally made - that the feedforward neural network could exploit similarity of positions. Although the setting of parameters in our initial runs involved some guesswork, now that we have a large set of "players" to examine, we can try to understand the phe- nomenon. Taking the champion networks at generation 1,000, 10,000, and 100,000 from our run, we sampled random players in their neighborhoods at different RMS distances to find out how likely is it to find a winning challenger. A thousand random neighbors at each of 11 different RMS distances played 8 games against the corresponding champion, and Figure 4 plots the fraction of games won by these challengers, as a function of RMS distance. This graph shows that as the players improve over time, the probability of finding good challengers in their neighborhood increases, which accounts for why the frequency of successful challenges goes up. 3 Each successive challenger is only required to 3. But why does the number of good challengers in a neighborhood go up, and if so, why does our algorithm falter nonetheless? There are several factors which require further study. It may be due to the general growth in weights, to less variability in strategy among mature players, or less ability simply to tell expert players apart with a few games. Figure 4: Distance versus probability of random challenger winning against champions at generation 1,000, 10,000 and 100,000. distance from champion 100k %wins for challenger take the small step of changing a few moves of the champion in order to beat it. The hope, for co-evolution, is that what was apparently unlearnable becomes learnable as we convert from a single question to a continuous stream of questions, each one dependent on the previous answer. 3.2. Replication Experiments After our first successful run, we tried to evolve ten more players using the same parameters and the same annealing schedule (10,000 and 70,000), but found that only one of these ten players was even competitive. Closer examination suggested that the other nine runs were failing because they were being annealed too early, before the frequency of successful challenges had reached an appropriate level. This premature annealing then made the task of the challengers even harder, so the challenger success rate fell even lower. We therefore abandoned the fixed annealing schedule and instead annealed whenever the challenger success rate exceeded 15% when averaged over 1000 generations. All ten players evolved under this regime were competitive (though not quite as good as our original player, which apparently benefitted from some extra inductive bias due to having its own tailor-made annealing schedule). Refining other heuristics and schedules could lead to superior players, but was not our goal. 3.3. Relative versus Absolute Expertise Does Backgammon allow relative expertise or is there some absolutely optimal strategy? Theoretically there exists a perfect "policy" for backgammon which would deliver the minimax optimal move for any position, and this perfect policy could exactly rate every other player on a linear scale, in practice, and especially without running 10,000 games to verify, it seems there are many relative cycles and these help prevent early convergence. In cellular studies of iterated prisoner's dilemma following (Axelrod, 1984) a stable population of "tit for tat" can be invaded by "all cooperate," which then allows exploitation by "all defect". This kind of relative-expertise dynamics, which can be seen clearly in the simple game of rock/paper/scissors (Littman, 1994) might initially seem very bad for self-play learning, because what looks like an advance might actually lead to a cycle of mediocrity. A small group of champions in a dominance circle can arise and hold a temporal oligopoly preventing further advance. On the other hand, it may be that such a basic form of instability prevents the formation of sub-optimal oligopolies and allows learning to progress. These problems are specific to non-zero-sum games; in zero sum games, appropriate use of self-play can be shown to converge to optimal play for both parties 4. Discussion We believe that our evidence of success in learning backgammon using simple hill-climbing in a relative fitness environment indicates that the reinforcement and temporal difference methodology used by Tesauro in his 1992 paper which led to TD-Gammon, while providing some advantage, was not essential for its success. Rather, a major contribution came from the co-evolutionary learning environment and the dynamics of back- gammon. Our result is thus similar to the bias found by Mitchell et al in Packard's evolution of cellular automata to the "edge of chaos" (Packard, 1988, Mitchell et al., 1993). Obviously, we are not suggesting that 1+1 hillclimbing is an advanced machine learning technique which others should bring to many tasks! Without internal cognition about an opponent's behavior, co-evolution usually requires a population. Therefore, there must be something about the domain itself which is helpful because it permitted both TD learning and hill-climbing to succeed through self-play, where they would clearly fail on other problem-solving tasks of this scale. In this section we discuss some issues about co-evolutionary learning and the dynamics of backgammon which may be critical to learning success. 4.1. Evolution versus Co-evolution TD-Gammon is a major milestone for a kind of evolutionary machine learning in which the initial specification of the model is far simpler than expected because the learning environment is specified implicitly, and emerges as a result of the co-evolution between a learning system and its training environment: the learner is embedded in an environment which responds to its own improvements - hopefully in a never-ending spiral, though this is an elusive goal to achieve in practice. While this co-evolutionary effect has been seen in population models, it is completely unexpected for a "1+1" hill-climbing evolution. Co-evolution has been explored on the sorting network problem (Hillis, 1992), on tic-tac-toe and other strategy games (Angeline and Pollack, 1994, Rosin and Belew, 1995, Schraudolph et al., 1994), on predator/prey games (Cliff and Miller, 1995, Reynolds, 1994) and on classification problems such as the intertwined spirals problem (Juille and Pollack, 1995). However, besides Tesauro's TD-Gammon, which has not to date been viewed as an instance of co-evolutionary learning, Sims' artificial robot game (Sims, 1994) is the only other domain as complex as backgammon to have had substantial success. Since a weak player can sometimes defeat a strong one, it should in theory be possible for a network to learn backgammon in a static evolutionary environment (playing against a fixed opponent) rather than a co-evolutionary one (playing against itself). Of course this is not as interesting an acheivement as learning without an expert on hand, and if TD-gammon had simply learned from Neurogammon, it wouldn't have been as startling a result. In order to further isolate the contribution of co-evolutionary learning, we had to modify our training setup because our original algorithm was only appropriate to self-play. In this new setup the current champion and mutant both play a number of games against the same opponent (called the foil) with the same dice-streams, and the weights are adjusted only if the champion loses all of these games while the mutant wins all of them. The number of pairs of games was initially set to 1 and incremented whenever the challenger success rate exceeded 15% when averaged over 1000 generations. The lower three plots in Figure 5, which track the performance of this algorithm with each of the three benchmark networks from our original experiments acting as foil, seem to show a relationship between learning rate and probability of winning. Against a weak foil (1k) learning is fast initially, when the probability of winning is around 50%, then tapers off as this probability increases. Against a strong foil (100k) learning is very slow initially, when the probability of winning is small, but speeds up as it increases towards 50%. All of these evolutionary runs were outperformed by a co-evolutionary version of the foil algorithm (co-ev) in which the champion network itself plays the role of the foil. Co-evolution seems to maintain a high learning rate throughout the run by automatically providing, for each new generation player, an opponent of the appropriate skill level to keep the probability of winning near 50%. Moreover, weaknesses in the foil are less likely to bias the learning process because they can be automatically corrected as the co-evolution proceeds (see also Section 4.3). 4.2. The Dynamics of Backgammon In general, the problem with learning through self-play discovered repeatedly in early AI and ML is that the learner could keep playing the same kinds of games over and over, only exploring some narrow region of the strategy space, missing out on critical areas of the game where it would then be vulnerable to other programs or human experts. This problem is particularly prevalent in deterministic games such as chess or tic-tac-toe. Tesauro (1992) pointed out some of the features of backgammon that make it suitable for approaches involving self-play and random initial conditions. Unlike chess, a draw is impossible and a game played by an untrained network making random moves will eventually terminate (though it may take much longer than a game between competent players). Moreover the randomness of the dice rolls leads self-play into a much larger part of the search space than it would be likely to explore in a deterministic game. We have worked on using a population to get around the limitations of self-play (Angeline and Pollack, 1994). Schraudolph et al., 1994 added non-determinism to the game of Go by choosing moves according to the Boltzmann distribution of statistical mechanics. Others, such as Fogel, 1993, expanded exploration by forcing initial moves. Epstein, 1994, has studied a mix of training using self-play, random testing, and playing against an expert in order to better understand these aspects of game learning. Generation Figure 5: Performance against PUBEVAL of players evolved by playing benchmark networks from our original run at generation 1k, 10k and 100k, compared with a co-evolutionary variant of the same algorithm. Each of these plots is an average over four runs. The performance of our original algorithm is included for comparison. original co-ev 100k We believe it is not enough to add randomness to a game or to force exploration through alternative training paradigms. There is something critical about the dynamics of backgammon which sets its apart from other games with random elements like Monopoly - namely, that the outcome of the game continues to be uncertain until all contact is broken and one side has a clear advantage. In Monopoly, an early advantage in purchasing properties leads to accumulating returns. What many observers find exciting about backgammon, and what helps a novice sometimes overcome an expert, is the number of situations where one dice roll, or an improbable sequence, can dramatically reverse which player is expected to win. In order to quantify this "reversibility" effect we collected some statistics from games played by our 100,000th generation network against itself. For each n between 0 and 120 we collected 100 different games in which there was still contact at move n, and, for n>6, 100 other games which had reached the racing stage by move n (but were still in Move Number Standard Deviation 1200.5Move Number contact racing game over Probability Figure (a) Standard deviation in the probability of winning for contact positions and racing positions. contact racing Figure (b) Probability of a game still being in the contact or racing stage at move n. Figure 7: Smoothed distributions of the probability of winning as a function of move number, for contact positions (left) and racing positions (right). Density Density Probability of Winning Move Number Move Probability progress). We then estimated the probability of winning from each of these 100 positions by playing out 200 different dice-streams. Figure 6 shows the standard deviation of this probability (assuming a mean of 0.5) as a function of n, as well as the probability of a game still being in the contact or racing stage at move n. Figure 7 shows the distribution in the probability of winning, as a function of move number, symmetrized and smoothed out by convolution with a gaussian function. These data indicate that the probability of winning tends to hover near 50% in the early stages of the game, gradually moving out as play proceeds, but typically remaining within the range of about 15% to 85% as long as there is still contact, thus allowing a reasonable chance for a reversal. These numbers could be different for other players, less reversability for stronger players perhaps and more for weaker ones, but we believe the effect remains an integral part of the game dynamics regardless of expertise. Our conjecture is that these dynamics facilitate the learning process by providing, in almost every situation, a nontrivial chance of winning and a nontrivial chance of losing, therefore potential to learn from the consequences of the current move. This is in deep contrast to many other domains in which early blunders could lead to a hopeless situation from which learning is virtually impossible because the reward has already become effectively unat- tainable. It seems this feature of backgammon may also be shared by other tasks for which TD-learning has been successful (Zhang and Dietterich, 1996, Crites and Barto, 1996, Walker et al., 1994). 4.3. Avoiding Suboptimal Equilibria in the Meta-Game of Learning A learning system can be viewed as an interaction between teacher and student in which the teacher's goal is to expose the student's weaknesses and correct them, while the student's goal is to placate the teacher and avoid further correction. We can build a model of this teacher/student interaction as a formal game, which we will call the Meta-Game of Learning (MGL) to avoid confusion with the game being learned. In this meta-game, the teacher T presents the student S with a sequence of questions prompting responses R i from the student. (In the backgammon domain, all the questions and responses would be legal positions, rolls and moves). S and T each receive payoffs in the process, which they attempt to maximize through their choices of questions and answers, and their limited abilities at self-modification. We generally assume the goal of learning is to prepare the student for interaction with a complex environment E that will provide an objective measure of its perfor- mance. 4 E and T thus play similar roles but are not assumed to be identical. The question then is: Can we find a payoff matrix for S and T which will enable the performance of S to continually improve (as measured by E)? If the rewards for T are too closely correlated with those for S, T may be tempted to ask questions that are too easy. If they are anti-correlated (for example if T=E), the questions might be too difficult. In either case it will be hard for S to learn (see Section 4.1). 4. For a general theory of evolution or self-organization, E is not necessary. An attractive solution to this problem is to have two or more students play the role of teacher for each other, or indeed a single student act as its own teacher, thus providing itself with questions that are always at the appropriate level of difficulty. The dynamics of the MGL, under such a self-teaching or co-evolutionary situation, would hopefully lead to a continuing spiral of improvement but may instead get bogged down by antagonistic or collusive dynamics, depending on the payoff structure. In our hillclimbing setup we may think of the mutant (teacher) trying to gain advantage (adjustment in the weights) by exploiting weaknesses in the champion, while the champion (student) is trying to avoid such an adjustment by not allowing its weaknesses to be exploited. Since the student and teacher are of approximately equal ability, it is to the advantage of the student to narrow the scope of the search, thus limiting the domain within which the teacher is able to look for a weakness. In most games, such as chess or tic-tac-toe, the student could achieve this by aiming for a draw instead of a win, or by always playing a particular style of game. If draws are not allowed, the teacher and student may figure out some other way to collude with each other - for example, each "throwing" alternate games (Angeline, 1994) by making a suboptimal sequence of early moves. These effects in self-learning systems, which may appear as early convergence in evolutionary algorithms, narrowing of scope, drawing or other collusion between teacher and student, are in fact Nash equilibria in the MGL, which we call Mediocre Stable States. 5 Our hypothesis is that certain features of backgammon operate against the formation of mediocre stable states in the MGL: backgammon is ergodic in the sense that any position can be reached from any other position 6 by some sequence of moves, and the dice rolls apparently create enough randomness to prevent either player from following a strategy that narrows the scope of the game appreciably. Moreover, early suboptimal moves are unlikely to provide the opponent with an easy win (see Section 4.2), so collusion by the throwing of alternate games is prevented. Mediocre stable states can also arise in human education systems, for example when the student gets all the answers right and rewards the teacher with positive teaching evaluations for not asking harder questions. In further work, we hope to apply the same kind of MGL equilibrium analysis to issues in human education. 5. Conclusions TD-Gammon remains a tremendous success in Machine Learning, but the causes for its success have not been well understood. The fundamental research in Tesauro's 1992 paper which was the basis for TD-Gammon, reportedly beat Sun's Gammontool 60- 65% of the time (depending on number of hidden units) and achieved parity against Neurogammon 1.0. Following this seminal 1992 paper, Tesauro incorporated a number of hand-crafted expert-knowledge features, eventually engineering a network which achieved world 5. MSS follows Maynard Smith's ESS (Maynard Smith, 1982) 6. with the exception of racing situations and positions with some pieces out of play. master level play (Tesauro, 1995). These features included concepts like existence of a prime, probability of blots being hit, and probability of escape from behind the oppo- nent's barrier. The evaluation function was also improved using multiple ply search. The best players we've been able to evolve can win about 45% of the time against PUBEVAL, which we believed to be about the same level at Tesauro's 1992 networks. Because Tesauro had never compared his 1992 networks to PUBEVAL, and because he used Gammontool's heuristic endgame in the ratings, the level of play achieved by these players has been somewhat murky: "the testing procedure is to play out the game with the network until it becomes a race, and then use Gammontool's algorithm to move for both sides until the end. This also does not penalize the TD net for having learned rather poorly the racing phase of the game."(p 272) When we compare our network's performance to PUBEVAL, it must be noted that we use our network's own (weak) endgame, rather than substituting in a much stronger expert system like Gammontool. Gerald Tesauro, in a commentary in this issue, has graciously cleared up the matter of comparing PUBEVAL to his 1992 results, and differs somewhat from our conclusions below. There are two other phenomena fom the 1992 paper, which are relevant to our work: "Performance on the 248-position racing test set reached about 65%. (This is substantially worse than the racing specialists described in the previous section.)" (p. "The training times .were on the order of 50,000 training games for the networks with games for the 20-hidden unit net, and 200,000 games for the 40-hidden unit net." (p. 273) Because we achieve similar levels of skills, and observe these same phenomena in training, endgame weakness, and convergence, we believe we have achieved results substantially similar to Tesauro's 1992 result, without any advanced learning algorithms. We could make stronger players by tuning the learning parameters, and adding more input features, but that is not our point. We do not claim that our 100,000th generation player is anywhere near as good as the current enhanced versions of TD-Gammon, ready to challenge the best humans, but it is surprisingly good considering its humble origins from hill-climbing with a relative fitness measure. Tuning our parameters or adding more input features would make more powerful players, but that was not the point of this study. We also do not claim there is anything "wrong" with TD learning, or that hillclimbing is just as good as reinforcement learning in general! Of course it isn't! Our point is that once an environment and representation have been refined to work well with a machine learning method, it should be benchmarked against the weakest possible algorithm so that credit for learning power can be properly distributed. We have noticed several weaknesses in our player that stem from the training which does not yet reward or punish the double and triple costs associated with severe losses ("gammoning" and "backgammoning") nor take into account the gambling process of "doubling." We are continuing to develop the player to be sensitive to these issues in the game. Interested players can challenge our 100,000th network using a web browser through our home page at: In conclusion, replicating some of Tesauro's 1992 TD-Gammon success under a much simpler learning paradigm, we find that the reinforcement and temporal difference methods are not the primary cause for success; rather it is the dynamics of backgammon combined with the power of co-evolutionary learning. If we can isolate the features of the backgammon domain which enable co-evolutionary and reinforcement learning to work so well, it may lead to a better understanding of the conditions neces- sary, in general, for complex self-organization. Acknowledgments This work was supported by ONR grant N00014-96-1-0418 and a Krasnow Foundation Postdoctoral fellowship. Thanks to Gerry Tesauro for providing PUBEVAL and subsequent means to calibrate it, Jack Laurence and Pablo Funes for development of the WWW front end to our evolved player, and comments from the Brandeis DEMO group, the anonymous referees, Justin Boyan, Tom Dietterich, Leslie Kaelbling, Brendan Kitts, Michael Littman, Andrew Moore, Rich Sutton and Wei Zhang. --R Competitive environments evolve better solutions for complex tasks. An alternate interpretation of the iterated prisoner's dilemma and the evolution of non-mutual cooperation The evolution of cooperation. Modular neural networks for learning context-dependent game strategies Tracking the red queen: Measurements of adaptive progress in co-evolutionary simulations Improving elevator performance using reinforcement learning. Massively parallel genetic programming. Markov games as a framework for multi-agent reinforcement learning Algorithms for Sequential Decision Making. Revisiting the edge of chaos: Evolving cellular automata to perform computations. Adaptation towards the edge of chaos. some studies of machine learning using the game of checkers. Temporal difference learning of position evaluation in the game of go. Evolving 3d morphology and behavior by competition. Learning to predict by the methods of temporal differences. Connectionist learning of expert preferences by comparison training. Practical issues in temporal difference learning. Temporal difference learning and td-gammon Temporal difference --TR --CTR Gerald Tesauro, Comments on Co-Evolution in the Successful Learning of Backgammon Strategy, Machine Learning, v.32 n.3, p.241-243, Sept. 1998 David B. Fogel, Beyond Samuel: evolving a nearly expert checkers player, Advances in evolutionary computing: theory and applications, Springer-Verlag New York, Inc., New York, NY, Ji Grim , Petr Somol , Pavel Pudil, Probabilistic neural network playing and learning tic-tac-toe, Pattern Recognition Letters, v.26 n.12, p.1866-1873, September 2005 multi-agent system integrating reinforcement learning, bidding and genetic algorithms, Web Intelligence and Agent System, v.1 n.3,4, p.187-202, December multi-agent system integrating reinforcement learning, bidding and genetic algorithms, Web Intelligence and Agent System, v.1 n.3-4, p.187-202, March Gerald Tesauro, Programming backammon using self-teaching neural nets, Artificial Intelligence, v.134 n.1-2, p.181-199, January 2002 Elizabeth Sklar , Mathew Davies, Multiagent simulation of learning environments, Proceedings of the fourth international joint conference on Autonomous agents and multiagent systems, July 25-29, 2005, The Netherlands Yeo Keun Kim , Jae Yun Kim , Yeongho Kim, A Tournament-Based Competitive Coevolutionary Algorithm, Applied Intelligence, v.20 n.3, p.267-281, May-June 2004 Elizabeth Sklar , Mathew Davies , Min San Tan Co, SimEd: Simulating Education as a Multi Agent System, Proceedings of the Third International Joint Conference on Autonomous Agents and Multiagent Systems, p.998-1005, July 19-23, 2004, New York, New York Edwin de Jong, The MaxSolve algorithm for coevolution, Proceedings of the 2005 conference on Genetic and evolutionary computation, June 25-29, 2005, Washington DC, USA Jordan B. Pollack , Hod Lipson , Gregory Hornby , Pablo Funes, Three generations of automatically designed robots, Artificial Life, v.7 n.3, p.215-223, Summer 2001 Pablo Funes , Jordan Pollack, Evolutionary Body Building: Adaptive Physical Designs for Robots, Artificial Life, v.4 n.4, p.337-357, October 1998 Frans A. Oliehoek , Edwin D. de Jong , Nikos Vlassis, The parallel Nash Memory for asymmetric games, Proceedings of the 8th annual conference on Genetic and evolutionary computation, July 08-12, 2006, Seattle, Washington, USA Jordan B. Pollack , Hod Lipson , Sevan Ficici , Pablo Funes , Greg Hornby , Richard A. Watson, Evolutionary techniques in physical robotics, Creative evolutionary systems, Morgan Kaufmann Publishers Inc., San Francisco, CA, 2001 Edwin D. de Jong, A Monotonic Archive for Pareto-Coevolution, Evolutionary Computation, v.15 n.1, p.61-93, Spring 2007 John Cartlidge , Seth Bullock, Combating Coevolutionary Disengagement by Reducing Parasite Virulence, Evolutionary Computation, v.12 n.2, p.193-222, June 2004 Stephan K. Chalup , Alan D. Blair, Incremental training of first order recurrent neural networks to predict a context-sensitive language, Neural Networks, v.16 n.7, p.955-972, September Edwin D. De Jong , Jordan B. Pollack, Ideal Evaluation from Coevolution, Evolutionary Computation, v.12 n.2, p.159-192, June 2004 Cooperative Multi-Agent Learning: The State of the Art, Autonomous Agents and Multi-Agent Systems, v.11 n.3, p.387-434, November 2005 Darse Billings , Lourdes Pea , Jonathan Schaeffer , Duane Szafron, Learning to play strong poker, Machines that learn to play games, Nova Science Publishers, Inc., Commack, NY, 2001

coevolution;reinforcement;temporal difference learning;self-learning;backgammon

298803

Locality Analysis for Parallel C Programs.

AbstractMany parallel architectures support a memory model where some memory accesses are local and, thus, inexpensive, while other memory accesses are remote and potentially quite expensive. In the case of memory references via pointers, it is often difficult to determine if the memory reference is guaranteed to be local and, thus, can be handled via an inexpensive memory operation. Determining which memory accesses are local can be done by the programmer, the compiler, or a combination of both. The overall goal is to minimize the work required by the programmer and have the compiler automate the process as much as possible. This paper reports on compiler techniques for determining when indirect memory references are local. The locality analysis has been implemented for a parallel dialect of C called EARTH-C, and it uses an algorithm inspired by type inference algorithms for fast points-to analysis. The algorithm statically estimates when an indirect reference via a pointer can be safely assumed to be a local access. The locality inference algorithm is also used to guide the automatic specialization of functions in order to take advantage of locality specific to particular calling contexts. In addition to these purely static techniques, we also suggest fine-grain and coarse-grain dynamic techniques. In this case, dynamic locality checks are inserted into the program and specialized code for the local case is inserted. In the fine-grain case, the checks are put around single memory references, while in the coarse-grain case the checks are put around larger program segments. The static locality analysis and automatic specialization has been implemented in the EARTH-C compiler, which produces low-level threaded code for the EARTH multithreaded architecture. Experimental results are presented for a set of benchmarks that operate on irregular, dynamically allocated data structures. Overall, the techniques give moderate to significant speedups, with the combination of static and dynamic techniques giving the best performance overall.

Introduction One of the key problems in parallel processing is to provide a programming model that is simple for the pro- grammer. One would like to give the programmer a familiar programming language, and have the programmer focus on high-level aspects such as coarse-grain parallelism and perhaps some sort of static or dynamic data distribution. Compiler techniques are then required to effectively map the high-level programs to actual parallel architectures. In this paper we present some compiler techniques that simplify the programmer's job when expressing locality of pointer data structures. As reported previously, we have developed a high-level parallel language called EARTH-C [1], and an associated compiler that translates EARTH-C programs to low-level threaded programs that execute on the EARTH multithreaded architecture [2, 3]. Our main emphasis is on the effective compilation of programs that use irreg- ular, dynamically-allocated data structures. Our initial approach provided high-level parallel constructs, and type extensions to express locality. The compiler then used the type declarations and dependence analysis to automatically produce low-level threads. Although our initial approach did provide a good, high-level, basis for programming the EARTH multi-threaded architecture, we found that the programmer was forced to make many function specializations, and to declare the appropriate pointer parameters and locally- scoped pointer variables as local pointers. Thus, in order to experiment with various locality approaches, the programmer needed to edit many places in his/her pro- gram, and to make several copies of the same function, each copy specialized for a particular type of locality. In order to ease the burden on the programmer we have developed some new compiler techniques to infer the locality of pointer variables, and then automatically produce the specialized versions of the functions. This allows the programmer to make very minimal changes to his/her high-level program in order to try various approaches to a problem. It also leads to shorter source programs because the programmer does not need to make several similar copies of the same function. The main idea behind our approach is that we use the information about the context of function calls and memory allocation statements to infer when indirect memory references must refer to local memory. We then automatically create specializations of functions with the appropriate parameters and locally-scoped variables explicitly declared as local pointers. This information is then used by the thread generator to reduce the number of remote operations required in the low-level threads. In order to test our approach we implemented the techniques in the EARTH-McCAT C compiler, and we experimented with a collection of pointer-based bench- marks. We present experimental measurements on the EARTH-MANNA machine to compare the performance of benchmarks without locality analysis, with locality analysis, and hand-coded versions with "the best" locality The rest of the paper is organized as follows. Section presents an overview of the EARTH-C language, the EARTH-McCAT compiler, and the EARTH-MANNA architecture. Section 3 provides some examples to motivate the locality analysis, and Section 4 describes the analysis itself. In Section 5 we give experimental results for our set of benchmarks programs. Finally, in Section 6 we discuss some related work, and in Section 7 we give conclusions and some suggestions on further work. 2 The EARTH-C Language The EARTH-C compiler has been designed to accept a high-level parallel C language called EARTH-C, and to produce a low-level threaded-C program that can be executed on the EARTH-MANNA multithreaded archi- tecture. In this section we provide an overview of the important points about the language, and the target ar- chitecture. More complete descriptions of the EARTH project can be found elsewhere [2]. 2.1 The EARTH-C Language The EARTH-C language has been designed with simple extensions to C. These extensions can be used to express parallelism via parallel statement sequences and a general type of forall loop; to express concurrent access via shared variables; and to express data locality via data declarations of local pointers. Any C program is a valid EARTH-C program, and the compiler will automatically produce a correct low-level threaded program. However, usually the programmer will make some minimal modifications to the program to expose coarse-grain parallelism, and to add information about data locality. Figure 1 gives two sample list processing functions, written in EARTH-C. In both cases the functions take a pointer to a list head, and a pointer to a node x, and return the number of times x occurs in the list. Figure 1(a) uses a forall loop to indicate that all interactions of the loop body may be performed in parallel. Since a loop must not have any loop-carried dependences on ordinary variables, we have used the shared variable count to accumulate the counts. Shared variables must always be accessed via atomic functions and in this case we have used the built-in functions writeto, addto and valueof. Figure 1(b) presents an alternative solution using recursion. In this example we use a parallel sequence (denoted using -" . "), to indicate that the call to equal node and the recursive call to count rec can be performed in parallel. The EARTH-C compiler captures coarse-grain parallelism at the level of function invocations. Ordinary C function calls are translated into lower-level threaded- C TOKEN calls. Such TOKEN calls are handled by the EARTH runtime load balancer, and the call will be mapped to a processor at runtime. However, in EARTH- C, it is also possible for the programmer to explicitly specify where the invocation should be executed using the syntax p(.)@expr. In this case the underlying threaded-C INVOKE mechanism is used to explicitly map the invocation to the processor specified by expr. For example, in Figure 1(b), the call to equal node is mapped to the processor owning node x, whereas the recursive call to count rec is not explicitly mapped to a processor, and it will be assigned by the runtime load balancer. In both cases, using the TOKEN and INVOKE mecha- nisms, the activation frame is allocated on the processor assigned the invocation, and the invocation will remain on the same processor for the lifetime of the invoca- tion. Thus, the EARTH-C compiler can assume that all parameters and locally-scoped variables are local memory accesses. On the contrary, since the invocations are mapped at runtime (either using the runtime load bal- ancer, or according to an expression that is evaluated at runtime), the compiler must assume that all accesses to global variables and all memory accesses via pointer in- directions, are to remote memory. Using these assumptions for function count rec in Figure 1(b), we can see that accesses to head, x, c1, and c2 are local accesses, but the access to head-?next is a remote memory ac- cess. Note that to make the locality easier to see, we underline all remote memory accesses, and we put local pointer declarations in bold type. As the target architecture for EARTH is a distributed- memorymachine, this distinction between local memory accesses and remote memory accesses is very important. Local memory accesses are expressed in the generated lower-level threaded-C program as ordinary C variables that are handled efficiently, and they may be assigned to registers or stored in the local data cache. However, remote memory references must be resolved by calls to the underlying EARTH runtime system. Thus, for remote memory accesses, there is the additional cost of the call to the appropriate EARTH primitive operation, plus the cost of accessing the communication network. If the remote memory access turns out to be actually on the same processor as the request, the communication time will be minimal, but it is still significantly more expensive than making a direct local memory access. Even though multithreaded architectures can hide some communication costs, it is clearly advantageous to maximize the use of local memory, whenever possi- ble. In order to expose more locality to the compiler, EARTH-C has the concept of local pointers. If the programmer knows that a pointer always points to local memory, then the keyword local may be added to the pointer type declaration. In Figure 1(a), all calls to equal node were made to the owner of the first argu- ment. Thus, in the declaration of equal node we have declared the first parameter to have type node local *p, which reading right to left, says that p is a pointer to a local node. Thus, in the body of equal node, the EARTH compiler may assume that p-?value is a local memory reference, but q-?value is potentially a remote memory reference. Figure 1(b) illustrates the opposite case, where the second parameter of equal node is a local pointer, and in this case the EARTH compiler must assume that p-?value is potentially remote, whereas q-?value is local. EARTH-C also includes another form of function declaration that also expresses locality. Functions may be declared using the keyword basic. These basic functions must only reference local memory, and they may not call any ordinary (remote) functions. Basic functions are translated into very cheap function invocations in the target threaded-C code, and all memory references within their bodies are ordinary C variable references. Thus, sometimes programmers use basic functions to indicate locality for all variable references int count(node *head, node *x) shared int count; node *p; if (equalnode(p,x)@OWNEROF(p)) int equalnode(node local *p, node *q) return(p-?value == q-?value); int countrec(node *head, node *x); node *next; int c1, c2; if (head != NULL) else int equalnode(node *p, node local *q) (a) iterative solution (b) recursive solution Figure 1: Example functions written in EARTH-C within a function body. The purpose of this paper is to help automate the generation of the local pointer declarations and to automatically provide specialized versions of functions for different calling contexts. Thus, the programmer concentrates on expressing where the computation should be mapped, and the compiler infers the locality information for pointers, and inserts the correct local pointer declarations. This reduces the burden on the program- mer, leads to shorter source programs, and makes changing the source program less error prone. In the examples in Figure 1, the programmer would only need to declare one version of equal node, and the compiler would automatically generate the appropriate specializations depending on the calling context. 2.2 The EARTH-McCAT C Compiler This paper builds upon the existing EARTH-McCAT C compiler. The overall structure of the compiler is given in Figure 2. The compiler is split into three phases. Phase I contains our standard transformations and anal- yses. The important points are that the source program is simplified into an AST-based SIMPLE intermediate representation [4]. At this point programs have been made structured via goto-elimination, and each statement has been simplified into a series of simple, basic statements. For each statement (including assignment statements, conditionals, loops, and function calls), we have the results of side-effect analysis that gives that set of locations read/written by the statement. The availability of this read/write information allows our locality analysis to be simple, and efficient. The methods presented in this paper are found in Phase II, where parallelization and locality enhancement is done. In these phases we use the results of the analyses from Phase I in order to transform the SIMPLE program representation into a semantically- equivalent program. The transformations presented in this paper introduce locality declarations, and produce new specialized versions of some functions. Phase III takes the transformed SIMPLE program from Phase II, generates threads, and produces the tar- Simplify Goto-Elimination Local Function Inlining Heap Analysis R/W Set Analysis Array Dependence Tester PHASE I Analyses and Transformations (Parallelization and Function Specialization Loop Partitioning Locality Analysis Points-to Analysis Thread Generation Build Hierarchical DDG Code Generation Locality Enhancement) Figure 2: Overall structure of the compiler get threaded-C code. By exposing more locality in Phase II, we allow the thread generator to deal with fewer remote memory accesses, and this should lead to fewer threads, fewer calls to EARTH primitives, and more efficient parallel programs. 2.3 The EARTH-MANNA Architecture In the EARTH model, a multiprocessor consists of multiple EARTH nodes and an interconnection network [2, 3]. As illustrated in Figure 3, each EARTH node consists of an Execution Unit (EU) and a Synchronization Unit (SU), linked together by buffers. The SU and EU share a local memory, which is part of a distributed shared memory architecture in which the aggregate of the local memories of all the nodes represents a global memory address space. The EU processes instructions in an active thread, where an active thread is initiated for execution when the EU fetches its thread id from the ready queue. The EU executes a thread to completion before moving to another thread. It interacts with the SU and the network by placing messages in the event queue. The SU fetches these messages, plus messages coming from remote processors via the network. The SU responds to remote synchronization commands and requests for data, and also determines which threads are to be run and adds their thread ids to the ready queue. EU SU EU SU Network Figure 3: The EARTH architecture Our experiments have been performed on a multi-threaded emulator built on top of the MANNA parallel machine[5]. Each MANNA node consists of two Intel CPUs, clocked at 50MHz, 32MB of dynamic RAM and a bidirectional network interface capable of transferring 50MB/S in each direction. The two processors on each node are mapped to the EARTH EU and SU. The EARTH runtime system supports efficient remote operations. Sequentially, loading a remote word takes about 7s, calling a remote function can be performed in 9s, and spawning a new remote thread takes about 4s. When issued in a pipeline these operation take only one third of these times. Motivating Examples In the preceding section (Figure 1) we presented an example where locality analysis could be used to make specialized versions of the equal node function. In this section we present some more typical examples of where locality information is used in EARTH-C programs, and we show how locality analysis and specialization can lead to better programs. These examples should give the intuitive ideas behind the actual locality analysis as presented in Section 4. In each of the example programs, all remote variable references are underlined. Thus, a program with fewer underlined references exhibits more locality, and will be more efficient. 3.1 Pointers to Local Variables and Parameters As outlined in Section 2.1, the underlying EARTH run-time system maps a function's activation frame to the processor executing the invocation, and thus it is safe to assume that parameters and locally-scoped variables are references to local memory. This assumption can be extended to pointer variables, if it can be shown that the pointer must point to locally-scoped variables and/or parameters. Figure 4(a) gives a somewhat contrived example that serves to illustrate the basic point. In function foo, pointer p points-to x, and x is a parame- ter. Since all parameters are allocated in local memory, it is safe to assume that *p points to local memory. Pointer q points either to parameter x, or to locally- scoped variable y. Since both x and y are local, we can assume that *q is local as well. Figure 4(b) gives the localized version of function foo. Note that the indirections *p and *q are remote (underlined) references in the original version of foo, but are local references in the localized version. int foo(int x) int y, *p, int *q; if (expr()) else int foo(int x) int y, local *p; int local *q; if (expr()) else (a) no locality inference (b) after locality inference Figure 4: Locality for Pointers 3.2 Dynamic Memory Allocation A rich source of locality information comes from the fact that dynamic memory allocation always allocates memory on the processor from which the allocator is called. Thus, if function f calls a memory allocation function like malloc, then the memory returned by malloc is local within the body of f. Consider the example in Figure 5(a). Without any locality inference, or type declarations, the compiler must assume that pointer t may refer to remote memory. Thus, as indicated by the underlined sections, all indirect references via t must be assumed to be possibly remote. However, one can note that t only points to the memory returned from malloc, and thus t can safely be declared as a local pointer, as illustrated in Figure 5(b). In this case, all memory accesses in the body of alloc point can be assumed to be local. 3.3 Mapping Computation to the OWNER OF Data The most common kind of locality information comes from the programmer mapping function invocations to the owner of a piece of data, using the @OWNER OF ex- pression. A typical example is given in Figure 6(a). The function count equal recursively descends through a binary tree t, counting the number of nodes with value v. The first recursive call, to the left sub-tree, is not explicitly mapped to any particular processor, and so there is no locality information for it. However, the second recursive call is explicitly sent to the owner of the right sub-tree. This means that these invocations can assume that all references via pointer t are local. As illustrated in Figure 6(b), to express this properly, a specialized copy of count equal must be created (called node *allocpoint(double x, double y, int colour) node * t; node *allocpoint(double x, double y, int colour) node local *t; (a) no locality inference (b) after locality inference Figure 5: Locality for Dynamic Memory Allocation count equal spec in the example), and in that copy the parameter t is declared to be a local pointer, and thus all memory accesses in the body are local. 3.4 Mapping Computation to HOME Another common method for mapping computation to specific processors is to use function calls of the form f(.)@HOME. This indicates that f should be invoked on the processor executing the call. From a locality standpoint, this gives us two kinds of information. First, if f returns a pointer value that is local within f, it must also be local within the body of function calling f. Second, if an argument to f is local in the caller, then the corresponding parameter must be local in the body of f. Figure 7(a) gives an example, and Figure 7(b) gives the result of applying locality analysis. First note that we can infer that t is a local pointer in the function newnode using the ideas presented for dynamic allocation given in section 3.2. Thus, the two calls to newnode in f must return local pointers, and both p and q must be local pointers. Now, consider the call to lessthan in f. Since both arguments, p and q, are local point- ers, the corresponding formals, a and b, in the body of lessthan must also be local pointers. 4 Locality Analysis In the last section, we identified the language/program features which are the sources of locality information. In this section we present the complete algorithm and associated analysis rules for locality analysis. The overall algorithm is presented in Figure 8. It works iteratively in two inter-related intra- and inter-procedural steps. At the beginning of the analysis, all the functions in the program are considered as candidates for specialization and put in the set spclPool. Further, the locality attribute for all formal parameters and global variables is initialized to Remote, unless the programmer has given explicit local pointer declarations. For pointers explicitly declared as local in the program, the locality attribute is initialized to Local. After this ini- tialization, the analysis proceeds in the following two steps: Step I: This step individually analyzes each function in the pool of functions to be specialized (spclPool). It starts with the current locality attribute of variables, and propagates this information throughout the procedure using a flow-insensitive intraprocedural approach. The details of this step are given in Subsection 4.1. Step II: This step performs interprocedural propagation of locality, and procedure specialization when ap- plicable. It looks at each call site in the functions belonging to spclPool, which is called with either HOME or OWNER OF primitive. Based on the locality information at the call-site, it infers locality information for formal parameters of the callee function as illustrated in sections 3.3 and 3.4. If a specialized version of the callee function with this locality already exists, the call-site is modified to invoke this function instead. Otherwise, a newly specialized version of the callee function is created for the given call-site. The locality attributes of the parameters of the newly-created function are appropriately initialized. As call-sites within the specialized function can trigger further specializations, the newly-created function is put in spclPool. At the end of this step, if spclPool is non-empty we go back to the first step. Clearly this process will terminate because we have only a finite number of func- tions/parameters that can be specialized, and specializations always add locality information. In the actual implementation, we just create a new locality context to represent a specialized function, and we do not actually create the complete new function. The decision to actually create specialized functions is taken after the analysis, depending upon the benefit achievable from a particular specialization. The details of the specialization step are given in Subsection 4.2. 4.1 Intraprocedural Locality Propagation We perform intraprocedural propagation of locality information using type-inference techniques [6], which have been previously adapted to perform almost linear points-to analysis[7]. The basic idea of the type inference algorithm is to partition program variables into a set of equivalence classes. To achieve this classification, a merging-based approach is used. For example, a simple assignment y, leads to assignment of the same type class to variables x and y, or in more general terms merging of the current type classes of x and y. If one wants to collect points-to information instead, the assignment would lead to merging of the points-to classes of variables x and y, where the points-to class of a variable contains the set of locations it may point to at runtime. Fast union/find data struc- int countequal(tree *t, int v) int c1, c2, c3; else int countequal(tree *t, int v) int c1, c2, c3; else int countequalspec(tree local *t, int v) int c1, c2, c3; else (a) no locality inference (b) after locality inference and specialization Figure Locality generated using OWNER OF tures can be used to make merging fast. This is the technique used by Steensgaard [7]. To collect locality information, we enhance this technique by attaching an additional locality attribute with each points-to class. The locality attribute can have one of the three possible values: (i) ?: indicating that the locality information is not yet determined, (ii) Local: all locations are definitely allocated on local memory, and (iii) Remote: some locations may be allocated on remote memory. When two points-to classes are merged, the new locality attribute is obtained by merging the locality attributes of the two classes using the merge operator ./ defined as follows: Remote Remote Remote Remote Below we give a small program fragment, and the points-to and locality information obtained by the type-inference based algorithm for it. int *a, *b, *c; int x, y, z; else The points-to and locality information at different program points is as follows: After After y, Localg y, Localg After y, Localg y, Localg The locality attribute is Local for all three pointers as they contain addresses of local variables. One can also note that the information provided is flow- insensitive, and there is no kill information (otherwise x should not be in the points-to class of a or b after statement S2). Thus the final information after S3 is conservatively valid for the entire program fragment. Our locality analysis uses the above type-inference based algorithm in an intraprocedural setting. The focus of the analysis is on accurately computing locality attributes, and not on computing complete points-to information. Thus, our analysis does not account for points-to information that holds due to aliasing between parameters and globals. However, since we make the worst case assumption about the locality of parameters and globals, this loss of information does not affect the correctness of our technique. Further, we have found that this loss of information does not affect the quality of the locality information that we find. Thus, it appears that this very inexpensive intraprocedural propagation is a good choice. In the following subsections we provide detailed rules for the intraprocedural locality analysis. Our analysis is performed at the SIMPLE intermediate representation of the EARTH-McCAT compiler. The SIMPLE representation provides eight basic statements that can affect points-to/locality information. Below, we provide void f() node *p, *q; if (lessthan(p,q)@HOME) - . node *newnode(int val) node *t; int lessthan(node *a, node *b) void f() node local *p, local *q; if (lessthan(p,q)@HOME) - . node *newnode(int val) node local *t; int lessthan(node local *a, node local *b) (a) no locality inference (b) after locality inference Figure 7: Locality generated using HOME locality analysis rules for each statement. 4.1.1 Address Assignment points to variable y, so we merge the points-to class of x with the class to which variable y belongs. if else 4.1.2 Dynamic Allocation For each statement S containing a mal- loc() call (or a related memory allocation call), we create a new variable called HeapS, and also create a class for it. The locality attribute of this class is initialized as Local. This is done because the EARTH programming model, requires a malloc call to always be allocated memory from the local processor. After creating this new class, we merge it with the points-to class of variable x, giving the following rule: 4.1.3 Pointer Assignments The statements belonging to this category include y, y, and the rules for analyzing them are discussed below. For the statement the operands y and z on right-hand side. This is conservatively safe, and the result does not depend on the operation being performed, or the type of the operands. For the statement y, we need to follow an additional level of indirection on left-hand side. We need to know what x points-to to perform the appropriate mergings, i.e. find the points-to-class of the points-to- class of x. If such a class does not yet exist, we simply create such a class, which gets filled in as the analysis proceeds. The same argument applies to statement x *y with respect to its right-hand side. The following table summarizes these rules. 4.1.4 Function Calls function call can considerably affect locality information. By using pointer arguments and global variables, it can modify the locality attribute, of any set of points-to classes. To avoid always making worst-case assumptions at function calls, locality analysis uses the results of interprocedural read-write (mod/ref) analysis, which is computed by our read-write set analysis in Step I of the compiler (re- fer to Figure 2 in Section 2.2). Based on the read-write information we have two important cases: Case I: The function call does not write to any pointer variable visible in the caller (including globals). This guarantees that the call does not affect the points-to and hence also locality information in the caller. In this case the statement statement only affect the points-to relationship/locality attribute of variable x. The locality attribute of the points-to class of x is updated depending upon the locality attribute of the points-to class of return f (return f is a symbolic name which represents the value returned by a function f), and the optional @expr used for the call. If the function f is a basic function, or called @HOME, and f returns a pointer (return f) pointing to Local, i.e. Locality(PointsToClass(return f)) is Local, it implies that the location pointed to by return f resides fun locality need to analyze all functions initially initialize locality(prog); while (spclPool != of functions to be analyzed is non\Gammaempty propagate locality(spclPool); fun propagate locality(prog, foreach func in spclPool = intraprocedural propagation propagate intraprocedural locality(func); deletFromSplcPool(func, no need to be analyzed again foreach callSite in prog = interprocedural propagation newly specialized function is created addToSpclPool(callSite.func, spclPool);=\Lambdanew func needs to be analyzed = fun propagate intraprocedural locality(func) = foreach assignmentStmt in func locality analyze stmt(assignmentStmt, localitySet); foreach callStmt in locality analyze call(callStmt, localitySet); if (callSite.type == @HOME jj callSite.type == @OWNER OF) find which params will be Local in the callee function specialzed version already exists with same locality set if (specialFuncExists(callSite.func, return new function is created else new func for callSite locality for new return new specialized function is created foreach func in prog conservatively assume all parameters and globals point\Gammato remote memory Figure 8: Overall Algorithm for Locality Analysis on the same processor. In this case, the new locality attribute of points-to class of x is obtained by merging it with Locality(PointsToClass(return f)). Otherwise it is simply assigned Remote, as per the rule below. if (expr == HOME Or IsBasicFunc(f)) else Case II: The other alternative is that the function call possibly writes to pointer variables of the caller. In this case we make worst-case assumptions and set the locality attribute of points-to classes of all arguments as Remote. We do it recursively for points-to classes of variables reachable via indirection on arguments : i.e. we also set Locality(P ointsT oClass (arg)) as Remote. We do not need to consider globals here as we already initialize their locality attribute as Remote at the start of the analysis. 4.2 Specializing Functions After completion of Step II of the algorithm, we have computed a set of possible specializations, and the associated locality. We use static estimates of the number of remote accesses saved to decide which specialized functions to actually create. For a given locality context of a function, we compute the following data: (i) weight of the corresponding call-site that reflects its potential execution frequency, and (ii) count of the remote accesses that can be eliminated by creating the specialized function. A specialized function is created for the given locality context if we find its (Weight * Count) estimate greater than our threshold, which we set to 20 by default. To compute the weight for call-sites, we first initialize the weight of all call sites to one. For each loop or recursion cycle, in which a call-site is embedded, its weight is multiplied by ten. The count of remote accesses saved is similarly estimated. A simple remote access saved is counted as one, while a remote access saved inside a loop is counted as ten. Further if some call sites inside the specialized function can also be specialized, we also add the number of remote accesses saved from this chain-specialization to the count. 5 Experimental Results In order to evaluate our approach, we have experimented with five benchmarks from the Olden suite [8], described in Table 1. All benchmarks use dynamic data structures (trees and lists) except quicksort, which uses dynamically-allocated arrays. The benchmark suite is suitable to evaluate our locality analysis focused on pointers Benchmark Description Problem Size power Optimization Problem based on a variable k-nary tree 10,000 leaves perimeter Computes the perimeter of a quad-tree encoded raster image Maximum tree- depth 11 quicksort Parallel version of quicksort 256K integers tsp Find sub-optimal tour for traveling problem cites health Simulates the Colombian health-care system using a 4-way tree 6 levels and 100 iterations Table 1: Benchmark Programs For each benchmark we provide results for three ver- sions: a simple EARTH-C version, a localized EARTH- C version and an advanced EARTH-C version. Hence- forth, we will refer to them respectively as simple, localized and advanced versions. The simple version implements the benchmark with the best data distribution that we have discovered to date for these benchmarks, and exploits locality using the @OWNER OF and @HOME primitives. It uses neither local pointers nor basic functions for this purpose. However, it can use basic functions for performing computations. The localized version is the benchmark obtained by applying our locality analysis and subsequent function specialization on the simple version. This version tries to find as many local pointers as possible. The advanced version is the hand-coded benchmark where the user tries to optimally exploit locality using local pointers, basic functions and other possible tricks. This version is based on the best efforts of our group to produce good speedups. None of these advanced versions were implemented by the authors Note that all three versions use the same general dynamic data distribution. However, the generated low-level program exploits locality to different degrees. Stat- ically, we can divide memory references into those that must be local, and those that might be remote. Each reference that must be local is translated into an ordinary C variable reference, which may be allocated to registers by the target compiler, or may be cached by the architecture. Each reference that might be remote is translated into a call to the EARTH runtime system. These calls to the runtime system may be resolved into memory requests to local memory, or to remote mem- ory, depending on the calling context. If at runtime, it resolves to a local memory reference, then we call this a pseudo-remote memory access. Note that pseudo- remote references are much more expensive than local references, but not as expensive as real remote references that must read or write data over the communication network. Our locality analysis and specializations effectively introduce more static declarations for local pointer vari- ables, and thus at runtime we execute fewer pseudo- remote references and more local references. As explained in the previous sections, this is done by automatically introducing local pointer declarations, and by introducing specialized versions of functions that capture locality for specific calling contexts. In Table 2 we summarize the static effect of applying our locality analysis and specializations to the simple versions of our benchmarks. For each benchmark, the first two columns list the number of local pointer declarations introduced, and the number of function specializations made in producing the localized version of the benchmark. The third column gives the relative sizes (lines of code) of the simple and advanced versions. Note that simple versions are all shorter, and sometimes substantially shorter than the advanced, hand- specialized programs. Benchmark #locals #spcls size(Simple) power perimeter Table 2: Static Measurements In Table 3 we provide data on the actual execution time, in milliseconds, for all three versions of each benchmark. These experiments were performed on the EARTH-MANNA architecture as described in Section 2.3. The last two columns give the percentage speedup obtained over the simple version for the localized and advanced versions respectively. For example, the column labeled "Localized vs. Simple" reports (T simple \Gamma localized )=T simple 100, and the column labelled "Ad- vanced vs. Simple" reports (T simple \GammaT advanced )=T simple 100. We provide data for benchmark runs on 1, 8 and processors. In Table 4 we present the actual number of remote data accesses and remote calls performed by different versions of each benchmark. The last two columns of this table give the percentage reduction in the number of remote data accesses/remote calls over the simple version for the localized and advanced versions respectively. Benchmark Simple Localized Advanced Localized Advanced EARTH-C EARTH-C EARTH-C vs. Simple vs. Simple (msec) (msec) (msec) (% impr) (%impr) power 1 proc 67158.06 64659.42 63482.45 3.72 5.47 8 procs 9132.92 8846.54 8651.86 3.14 5.27 perimeter 1 proc 7095.55 5966.37 5255.03 15.91 25.90 8 procs 1220.71 894.86 872.59 26.70 28.50 procs 748.31 546.23 523.32 27.00 30.06 8 procs 5394.66 5020.13 4587.74 6.94 15.00 8 procs 17193.55 15104.78 1116.16 12.10 93.50 Table 3: Execution Time Note that the number of remote accesses performed is independent of the number of processors used for a given program, when we do not differentiate between a real remote access and a pseudo-remote access. The data in Tables 3 and 4 indicates that the localized version always performs better than the simple version, i.e. our locality analysis is always able to identify some additional locality. Further, the percentage improvement can vary a lot, depending upon the bench- mark. The localized version comes very close to the advanced version for the first three of the five bench- marks. In the last two cases the localized version does give an improvement, but does not compete with the hand-coded advanced version. We analyze these results in detail individually for each benchmark below. Power: This benchmark implements the power system optimization problem [9]. It uses a four-level tree structure with different branching widths at each level. Our locality analysis achieves 3-4% improvement over the simple version in execution time, which is quite close to the advanced version (5%). However, it is able to achieve an 80% reduction in the number of remote data accesses (Table 4). This happens because function calls in this benchmark are typically of the format compute node(node)@OWNER OF(node), and the function performs numerous scalar data accesses of the form (*node).item. Our analysis captures the locality of the pointer node and eliminates all remote data accesses with respect to it. The significant reduction in remote accesses does not reflect in the execution time, as the benchmark spends most of the time in performing floating point operations, which far exceeds the time spent in data accesses. The advanced version achieves 93% reduction in the number of remote calls over both the simple and localized version, by using basic functions. This factor enables it to achieve slightly better speedup than the localized version. Perimeter: This benchmark computes the perimeter of a quad-tree encoded raster image [9]. The unit square image is recursively divided into four quadrants until each one has only one point. The tree is then traversed bottom-up to compute the perimeter of of each quadrant The localized version achieves 15-27% speedup and comes very close to the advanced version for the 8 and processor runs. The localized version has 32% fewer remote accesses. This reduction more significantly affects the execution time than for power benchmark, because the benchmark does not involve much computation and spends most of its time in traversing the quad-tree and hence performing data accesses. It is an irregular benchmark, and each computation requires accesses to tree nodes which may not be physically close to each other. Due to this characteristic, the advanced version cannot exploit additional locality using basic functions. The localized version thus competes very well with it. Quicksort: This benchmark is a parallel version of the standard quicksort algorithm. The two recursive calls to quicksort are executed in parallel, with the call for the bigger subarray invoked @HOME. Because the size of subarrays in each recursive sorting phase is unknown in advance, dynamically-allocated arrays are used. In the simple version, function qsort copies the in-coming array to a local array using a blkmov, and at the end copies back the local array to the incoming array using another blkmov. Our locality analysis is able to identify the locality of the incoming array for the @HOME call. It generates a specialized version of the recursive qsort function, with the incoming array declared as local pointer, and the two blkmov instructions substituted by calls to the basic function memcpy. This transformation enables the localized version to achieve a significant 80% speedup over the simple version, which is within 1% of the speedup obtained by the advanced version. The advanced version uses some additional basic functions to completely eliminate remote calls, but performs just a little better than the localized version. Benchmark Simple Localized Advanced Localized Advanced EARTH-C EARTH-C EARTH-C vs. Simple vs. Simple power 2294179 451179 204179 80.33 91.10 perimeter 2421800 1635111 1586323 32.48 34.49 quicksort 8498635 29128 216 99.65 99.99 tsp 4421050 2672068 829790 39.56 81.23 health 41409726 33148575 606 19.94 99.99 Table 4: Remote Accesses Saved Tsp: This benchmark solves the traveling salesman problem using a divide-and-conquer approach based on close- point algorithm [9]. This algorithm first searches a sub-optimal tour for each subtree(region) and then merges subtours into bigger ones. The tour found is built as a circular linked list sitting on top of the root nodes of subtrees. Similar to perimeter, this benchmark is irregular in nature and spends a significant amount time on data accesses. The localized version achieves 6-8% speedup from a 39% reduction in remote data accesses. The localized version, however, fails to compete with the advanced version, which achieves upto 20% speedup resulting from an 81% reduction in remote data ac- cesses. This happens because the linked lists representing tours, are distributed in segments and there are only very few links across processors. With the knowledge that an entire sublist is local, the advanced version exploits significant data locality by using basic functions to traverse these local sublists. Our locality analysis is not designed to identify locality of recursive fields. This kind of locality is implicit in the programmer's organization of the data, and is very difficult to find with compiler analyses. This benchmark simulates the Colombian health-care system using a 4-way tree [9]. Each village has four child villages, and a village hospital, treating patients from the villages in the same subtree. At each time step, the tree is traversed, and patients, once assessed, are either treated or passed up to the parent tree node. The 4-way tree is evenly distributed among the processors and only top-level tree nodes have their children spread among different processors. With the locality analysis, we are able to achieve up to 12% speedup, resulting from a 19% reduction in remote data accesses. This benchmark is similar to power in its call pattern. However, for a call of the format foo(village node)@OWNER OF(village node), it recursive data structures through village node (linked list of patients) as opposed to scalar data items like in power. Our analysis is able to eliminate all remote accesses with respect to the village node like list (*village node).list, but not further accesses like (*list).patient. Thus the localized version gets decent speedup over the simple version, but does not compete with the advanced version. With the knowledge that only top level tree nodes can have remote children, the advanced version eliminates almost all the remote data accesses and calls. In this regard, this benchmark is similar to tsp. 5.1 Summary In summary, we find that our locality analysis does give significant improvements in all cases. However, in some cases locality analysis cannot compete with hand-coded locality mapping and declarations provided by the pro- grammer. As the programmer may have some implicit knowledge of data locality, it is important to retain the ability to explicitly declare local pointers. Since in many cases the declarations can be auto- mated, and the locality analysis and specialization is a source-to-source transformation, one could imagine that the programmer could use the output of the compiler to produce a localized version of the program, and then test the program to see if acceptable speedup is achieved. If there is more locality in the program, then the programmer could add further locality declarations in order to further improve the program. Finally, note that our experiments are performed on EARTH- MANNA. Compared with other distributed memory systems like the IBM SP2, or a network of workstations, EARTH-MANNA has a much smaller memory latency, which sometimes can be further hidden by multi-threading techniques. Therefore, we can expect better speedup from locality analysis for those machines with larger memory latencies and those that do not support multi-threading techniques. 6 Related Work Our intraprocedural locality propagation approach is similar to Steensgaard's [7] linear points-to analysis al- gorithm. Other related work in the area of parallelizing programs with dynamic data structures is that of Carlisle [9] for distributed memory machines, and of Rinard and Diniz [10] for shared memory machines. In particular, Carlisle's [9] affinity analysis has similar goals as our locality analysis, albeit we target different kinds of locality. Steensgaard proposed a type-inference based algorithm for points-to analysis with almost linear time com- plexity. His algorithm is both flow-insensitive and context- insensitive. On encountering a call-site he simply merges the formal parameters with their respective actual argu- ments. Thus his algorithm cannot distinguish between information arriving at a function from different calling contexts. For our locality analysis, the calling context information is crucial. The invocation specification (@HOME, @OWNER OF) of the call site is a major source of locality information. Thus we do not want to merge locality information arriving from different calling contexts. On the contrary, we want to create specialized versions of a function for each calling context that provides us substantial locality. To this end, we use type-inference to only propagate information intraprocedurally, and employ a different technique for interprocedural propagation as explained in section 4. However, to ensure that our intraprocedural propagation collects conservatively safe information, we need to make conservative assumptions at the start of the procedure and on encountering call-sites (using read-write set information). Thus, although our approach was originally inspired by the points-to analysis, it is really specifically tailored to capture the information most relevant to locality analysis Carlisle's affinity analysis is designed to exploit the locality with respect to linked fields. His analysis relies upon the information regarding the probability (affin- ity) that nodes accessible by traversing a linked field are residing on the same processor. If the affinity is high, then he puts runtime checks to eliminate remote accesses. However, the affinity information is not infered by the compiler, but is provided via programmer annotations. Our locality analysis is not designed to exploit the locality achievable via linked fields (as discussed in section because we wanted our analysis to be an automatic compiler analysis, and we did not want to burden the user for additional detailed information. Locality of recursive fields can be explicitly declared using local pointer declarations in EARTH-C. 7 Conclusions and Further Work In this paper we have presented a locality analysis for parallel C programs, based on type-inference techniques. Our analysis tries to exploit additional fine-grain locality from the coarse-grain locality information already provided by the user and from other program characteristics like malloc sites. We evaluated its effectiveness for a set of benchmarks and found that it can eliminate significant number of pseudo-remote accesses and provide speedups ranging from a modest 4% up to 80% over the original parallel program. For several benchmarks, this speedup also comes very close to the speedup obtained by an advanced hand-coded version. Further- more, we found that the locality analysis reduced the burden on the programmer, and allowed us to develop shorter, more general benchmarks. Based on the encouraging results from this paper, we plan to evaluate our analysis on a wider set of bench- marks. We also plan to use flow-sensitive locality propagation techniques, that can exploit the locality of a pointer, even if it is local within only a specific section of the program. Another goal is to automatically identify basic functions. Finally, as the locality information for linked fields can sometimes provide significant speedups (for benchmarks health and tsp), we plan to extend our analysis to capture this type of locality, using profile information, and efficiently-scheduled runtime checks. Acknowledgements We gratefully acknowledge the support from people in the EARTH group, specially Prof. Guang R. Gao and Olivier Maquelin. This research was funded in part by NSERC, Canada. --R "Compiling C for the EARTH multithreaded architecture," "A study of the EARTH-MANNA multithreaded system," "Polling Watchdog: Combining polling and interrupts for efficient message handling," "Designing the McCAT compiler based on a family of structured intermediate representations," "La- tency hiding in message-passing architectures," "Efficient type inference for higher-order binding-time analysis," "Points-to analysis in almost linear time," "Supporting dynamic data structures on distributed-memory machines," Olden: Parallelizing Programs with Dynamic Data Structures on Distributed-Memory Machines "Commutativity analysis: A new analysis framework for parallelizing compilers," --TR --CTR Francisco Corbera , Rafael Asenjo , Emilio L. Zapata, A Framework to Capture Dynamic Data Structures in Pointer-Based Codes, IEEE Transactions on Parallel and Distributed Systems, v.15 n.2, p.151-166, February 2004 Oscar Plata , Rafael Asenjo , Eladio Gutirrez , Francisco Corbera , Angeles Navarro , Emilio L. Zapata, On the parallelization of irregular and dynamic programs, Parallel Computing, v.31 n.6, p.544-562, June 2005

compiling for parallel architectures;locality analysis;multithreaded architectures

298815

An Index-Based Checkpointing Algorithm for Autonomous Distributed Systems.

AbstractThis paper presents an index-based checkpointing algorithm for distributed systems with the aim of reducing the total number of checkpoints while ensuring that each checkpoint belongs to at least one consistent global checkpoint (or recovery line). The algorithm is based on an equivalence relation defined between pairs of successive checkpoints of a process which allows us, in some cases, to advance the recovery line of the computation without forcing checkpoints in other processes. The algorithm is well-suited for autonomous and heterogeneous environments, where each process does not know any private information about other processes and private information of the same type of distinct processes is not related (e.g., clock granularity, local checkpointing strategy, etc.). We also present a simulation study which compares the checkpointing-recovery overhead of this algorithm to the ones of previous solutions.

Introduction Checkpointing is one of the techniques for providing fault-tolerance in distributed systems [6]. A global check-point consists of a set of local checkpoints, one for each pro- cess, from which a distributed computation can be restarted after a failure. A local checkpoint is a state of a process saved on stable storage. A global checkpoint is consistent if no local checkpoint in that set happens before [9] another one [4, 10]. Three classes of algorithms have been proposed in the literature to determine consistent global checkpoints: un- coordinated, coordinated and communication induced [6]. In the first class, processes take local checkpoints independently on each other and upon the occurrence of a failure, a procedure of rollback-recovery tries to build a consistent global checkpoint. Note that, a consistent global checkpoint might not exist producing a domino effect [1, 12] which, in the worst case, rollbacks the computation to its initial state. In the second class, an initiator process forces other pro- cesses, during a failure-free computation, to take a local This work is partially supported by Scientific Cooperation Net-work of the European Community "OLOS" under contract No. ERB4050PL932483. checkpoint by using control messages. The coordination can be either blocking [4] or non-blocking [8]. However, in both cases, the last local checkpoint of each process belongs to a consistent global checkpoint. In the third class, the coordination is done in a lazy fashion by piggybacking control information on application messages. Communication-induced checkpointing algorithms can be classified in two distinct categories: model-based and index-based [6]. Algorithms in the first cate- gory, for example [2, 14], have the target to mimic a piece-wise deterministic behavior for each process [7, 13] as well as providing the domino-free property. Index-based algorithms associate each local checkpoint with a sequence number and try to enforce consistency among local checkpoints with the same sequence number [3, 5, 11]. Index-based algorithms ensure domino-free rollback with, gener- ally, less overhead, in terms of number of checkpoints and control information, than model-based ones. In this paper we present an index-based checkpointing protocol that reduces the checkpointing overhead in terms of number of checkpoints compared to previous index-based algorithms. Our protocol is well suited for autonomous environments where each process does not have any private information of other processes. To design our algorithm, we extract the rules, used by index-based algorithms, to update the sequence number. This points out that checkpoints are due to the process of increasing of the sequence numbers. Hence, we derive an algorithm that, by using an equivalence relation between pair of successive checkpoints of a process, allows a recovery line to advance without increasing its sequence number. In the worst case, our algorithm takes the same number of checkpoints as in [11]. The advantages of our algorithm are quantified by a simulation study showing that the check-pointing overhead can be reduced up to 30% compared to the best previous solution. The price we pay is that each application message piggybacks more control information (one vector of integers) compared to previous proposals. The paper is organized as follows. Section 2 presents the system model. Section 3 shows the class of index-based checkpointing algorithms in the context of autonomous en- vironments. Section 4 describes the equivalence relation and the proposed algorithm. In Section 5 a simulation study is presented. 2. Model of the Distributed Computation We consider a distributed computation consisting of n which interact by messages ex- changing. Each pair of processes is connected by a two- ways reliable channel whose transmission delay is unpredictable but finite. are autonomous in the sense that: they do not share memory, do not share a common clock value 1 and do not have access to private information of other processes such as clock drift, clock granularity, clock precision and speed. Recovery actions due to process failures are not considered in this paper. A process produces a sequence of events; each event moves the process from one state to another. We assume events are produced by the execution of internal, send or receive statements. Moreover, for simplicity, we consider a checkpoint C as a particular type of internal event of a process, which dumps the current process state onto stable storage. The send and receive events of a message m are denoted respectively with send(m) and receive(m). A distributed execution - can be modeled as a partial order of events - is the set of all events and ! is the happened-before relation [9]. A checkpoint in process P i is denoted as C i;sn where sn is called the index, or sequence number, of a checkpoint. Each process takes checkpoints either at its own pace (ba- sic checkpoints) or induced by some communication pattern (forced checkpoints). We assume that each process P i takes an initial basic checkpoint C i;0 and that, for the sake of sim- plicity, basic checkpoints are taken by a periodic algorithm. We use the notation next(C i;sn ) to indicate the successive checkpoint, taken by P i , after C i;sn . A checkpoint interval I i;sn is the set of events between C i;sn and next(C i;sn ). message m sent by P i to P j is called orphan with respect to a pair (C i;sn its receive event occurred before C j;sn j while its send event occurred after checkpoint C is a set of local checkpoints one for each process. A global checkpoint C is consistent if no orphan message exists in any pair of local checkpoints belonging to C. In the fol- lowing, we denote with C sn a global checkpoint formed by checkpoints with sequence number sn 2 . 1 The index-based algorithm presented in [5] assumes, for example, a standard clock synchronization algorithm, which provides a common clock value to each process. We use the term consistent global checkpoint Csn and recovery line Lsn interchangeably. 3. Index-based Checkpointing Algorithms The simplest way to form a consistent global checkpoint is, each time a basic checkpoint C i;sn is taken by process to start an explicit coordination. This coordination results in a consistent global checkpoint C sn associated to that local checkpoint. This strategy induces checkpoints (one for each process) per basic checkpoint. Briatico at al. [3] argued that the previous "centralized" strategy can be "decentralized" in a lazy fashion by piggy-backing on each application message m the index sn of the last checkpoint taken (denoted m:sn). Let us assume each process P i endows a variable sn i which represents the sequence number of the last check- point. Then, the Briatico-Ciuffoletti-Simoncini (BCS) algorithm can be sketched by using the following rules associated with the action to take a local checkpoint: When a basic checkpoint is scheduled, sn i is increased by one and a checkpoint C i;sn i is taken; Upon the receipt of a message m, if sn checkpoint Ci;m:sn is taken and sn i is set to m:sn, then the message is processed. By using the above rules, it has been proved that C sn is consistent [3]. Note that, due to the rule take-forced(BCS), there could be some gap in the index assigned to checkpoints by a process. Hence, if a process has not assigned the index sn, the first local checkpoint of the process with sequence number greater than sn can be included in the consistent global checkpoint C sn . In the worst case of BCS algorithm, the number of forced checkpoints induced by a basic one is n \Gamma 1. In the best case, if all processes take a basic checkpoint at the same physical time, the number of forced checkpoints per basic one is zero. However, in an autonomous environment, local indices of processes may diverge due to many causes (process speed, different period of the basic checkpoint etc). This pushes the indices of some processes higher and each time one of such processes sends a message to another one, it is extremely likely that a number of forced checkpoints, close to will be induced. To reduce the number of checkpoints, an interesting observation comes from the Manivannan-Singhal algorithm [11] which has been designed for non-autonomous distributed systems. There is no reason to take a basic check-point if at least one forced checkpoint has been taken during the current checkpoint interval. So, let us assume process indicates if at least one forced checkpoint is taken in the current checkpoint period (this flag is set to FALSE each time a basic checkpoint is scheduled, and set to TRUE each time a forced checkpoint is taken). A version of Manivannan-Singhal (MS) algorithm well suited for autonomous environment can be sketched by the following rules: When a basic checkpoint is scheduled, if skip then increased by one and a checkpoint Upon the receipt of a message m, if sn i ! m:sn then a checkpoint Ci;m:sn is taken, sn i is set to m:sn and skip the message is processed. Even though MS algorithm produces a reduction of the total number of checkpoints, the number of forced checkpoints caused by a basic one is equal to BCS as take-forced(MS) is actually similar to take-forced(BCS). In this section we propose an algorithm that includes the take-basic(MS) rule, however, when a basic check-point is taken, the local sequence number is updated only if there was the occurrence of a particular checkpoint and communication pattern. The rationale behind this solution is that each time a basic checkpoint is taken without increasing the sequence number, it does not force checkpoints and this reduces the total number of checkpoints. At this end, let us first introduce a relation of equivalence defined on pairs of successive checkpoints of a process. 4.1. Equivalence Relation Between Checkpoints Definition 4.1 Two local checkpoints C i;sn and next(C i;sn ) of process P i are equivalent with respect to the recovery line L sn , denoted Lsn Lsn sn sn I 2;sn Figure 1. Examples of pairs of equivalent checkpoints. As an example, let consider the recovery line L sn depicted in Figure 1. If in I 2;sn process P 2 executes either send events or receive events of messages which have been sent before the checkpoints included in the recovery line Lsn recovery line L 0 sn is created by replacing C 2;sn with next(C 2;sn ) from L sn . Figure also shows the construction of the recovery line sn starting from L 0 sn by using the equivalence between C 1;sn and next(C 1;sn ) with respect to L 0 sn . As shown in the above examples, the equivalence relation has a simple property (see Lemmas 4.1 and 4.2 of Section Lsn then the set L 0 fC i;sn g[fnext(C i;sn )g is a recovery line. Hence, the presence of a pair of equivalent checkpoints allows a process to locally advance a recovery line without updating the sequence 4.2. Sequence and Equivalence Numbers of a Recovery line We suppose process P i owns two local variables: sn i (sequence number) and en i (equivalence number). The variable sn i stores the number of the current recovery line. The variable en i represents the number of equivalent local checkpoints with respect to the current recovery line (both sn i and en i are initialized to zero). Hence, we denote as C i;sn;en the checkpoint of P i with the sequence number sn and the equivalence number en; the en ? is also called the index of a check- point. Thus, the initial checkpoint of process P i will be denoted as C i;0;0 . The index of a checkpoint is updated according to the following rule: if C i;sn;en Lsn then next(C i;sn;en C i;sn+1;0 . Process P i also endows a vector EQ i of n integers. The j-th entry of the vector represents the knowledge of P i about the equivalence number of P j with the current sequence number sn i (thus the i-th entry corresponds to en i ). EQ i is updated according to the following rule: each application message m sent by process P i piggybacks the current sequence number sn i (m:sn) and the current EQ i vector (m:EQ). Upon the receipt of a message m, if is updated from m:EQ by taking a component-wise maximum. If m:sn ? sn i , the values in m:EQ and m:sn are copied in EQ i and sn i . Let us remark that the set L recovery line (a sketch proof of this property is given in Lemma 4.4). So, to the knowledge of P i , the vector EQ i actually represents the most recent recovery line with sequence number sn i . 4.3. Tracking the Equivalence Relation Upon the arrival of a message m, sent by P j , at P i in the checkpoint interval I i;sn;en , one of the following three cases is true: (m has been sent from the left side of the recovery line has been sent from the right side of the recovery line [ 8j C j;sn;EQ i [j] ); been sent from the right side of a recovery line of which P i was not aware). According to previous cases, at the time of the check-point next(C i;sn;en ), in one of the following three alternatives: (i) If no message m is received in I i;sn;en that falls in case 2 or 3, then C i;sn;en Lsn next(C i;sn;en ). That equivalence can be tracked by a process using its local context at the time of the checkpoint next(C i;sn;en ). Thus next(C i;sn;en equivalence between shown in Figure 2, is an example of such a behavior). (ii) If there exists a message m which falls in case 3 then C i;sn;en is not equivalent to next(C i;sn;en ) and thus next(C i;sn;en (iii) If no message falls in case 3 and there exists at least a message m received in I i;sn;en which falls in case 2, then the checkpoint next(C i;sn;en ) is causally related to one checkpoint belonging to the recovery line formed by communication pattern is shown in Figure where, due to m, C 2;sn;0 ! next(C 1;sn;0 )). The consequence is that process P i cannot determine, at the time of taking the checkpoint next(C i;sn;en ), if Lsn optimistically (and provisionally) that C i;sn;en Lsn As provisional indices cannot be propagated in the system, if at the time of the first send event after next(C i;sn;en ) the equivalence is still undetermined, then next(C i;sn;en 0). Otherwise, the provisional index becomes permanent. Figure 2 shows a case in which message m 0 brings the information (encoded in m 0 :EQ) to P 1 (before the sending of Lsn and the recovery line was advanced, by P 2 , from L sn to L 0 sn . In such a case, P 1 can determine C 1;sn;0 is equivalent to next(C 1;sn;0 ) with respect to L 0 sn and, then, advances the recovery line to L 00 sn . 4.4. The Algorithm The checkpointing algorithm we propose (BQF) takes basic checkpoints by using the take-basic(MS) rule. However, it does not update the sequence number by optimistically assuming that a basic checkpoint is equivalent to the previous one. So we have: Lsn sn sn Figure 2. Upon the receipt of m 0 , P 1 detects sn When a basic checkpoint is scheduled, If skip i then skip else en checkpoint Ci;sn;en is taken with index provisionally set to ! sn Due to the presence of provisional indices caused by the equivalence relation, our algorithm needs a rule, when sending a message, in order to disseminate only permanent indices of checkpoints. before sending a message m in I i;sn i ;en i , if there has been no send event in I i;sn i ;en i and the index is provisional then Lsn then the index ! sn permanent else sn and the index of the last checkpoint is replaced permanently with ! sn the message m is sent; The last rule of our algorithm take-forced(BQF) refines BCS's one by using a simple observation. Upon the receipt of a message m such that m:sn ? sn i , there is no reason to take a forced checkpoint if there has been no send event in the current checkpoint interval I i;sn;en . In- deed, no causal relation can be established between the last checkpoint C i;sn i ;en i and any checkpoint belonging to the recovery line L m:sn and, thus, the index of C i;sn i ;en i can be replaced permanently with the index ! m:sn; 0 ?. take-forced(BQF): Upon the receipt of a message m in I i;sn i ;en (a) If sn i ! m:sn and there has been at least a send event in I i;sn i ;en i then begin a forced checkpoint Ci;m:sn;0 is taken and its index is permanent; (b) If sn i ! m:sn and there has been no send event in I i;sn i ;en i then begin the index of the last checkpoint C i;sn i ;en i is replaced permanently the message m is processed; For example, in Figure 3.a, the local checkpoint C 3;sn;en3 can belong to the recovery line L sn+1 (so the index can be replaced with forced checkpoint (b) C3;sn;en 3 C3;sn+1;0 next(C2;sn;en C3;sn;en 3 (a) Figure 3. Upon the receipt of m, C 3;sn;en3 can be a part of L sn+1 (a); C 3;sn;en3 cannot belong to L sn+1 (b). the contrary, due to the send event in I 3;sn;en3 depicted in Figure 3.b, a forced checkpoint with index has to be taken before the processing of message m. Point (b) of take-forced(BQF) decreases the number of forced checkpoints compared to BCS. The else alternative of send-message(BQF) and the part (a) of take-forced(BQF), represent the cases in which the action to take a basic checkpoint leads to update the sequence number with the consequent induction of checkpoints in other processes. Data Structures and Process Behavior. We assume each process P i has the following data structures: after first send i , skip i , provisional past The boolean variable after first send i is set to TRUE if at least one send event has occurred in the current check-point interval. It is set to FALSE each time a checkpoint is taken. The boolean variable provisional i is set to TRUE whenever a provisional index assignement occurs. It is set to FALSE whenever the index becomes permanent. present i [j] represents the maximum equivalence number en j piggybacked on a message m received in the current checkpoint interval by P i and that falls in the case 2 of Section 4.3. Upon taking a checkpoint or when updating the sequence number, present i is initialized to -1. If the checkpoint is basic, present i is copied in past i before its initialization. Each time a message m is received such that past past i [j] is set to -1. So, the predicate past indicates that there is a message received in the past checkpoint interval that has been sent from the right side of the recovery line (case 2 of Section currently seen by P i . Below the process behavior is shown (the procedures and the message handler are executed in atomic fashion). This implementation assumes that there exist at most one provisional index in each process. So each time two successive provisional indices are detected, the first index is permanently replaced with ! sn init en i := 0; after first send i := FALSE; past i [h] := \Gamma1; 8h present i [h] := \Gamma1; when (m) arrives at P i from begin if after first send i then begin take a checkpoint C; % forced checkpoint % after first send i := FALSE; assign the index ! sn to the last checkpoint C; provisional i := FALSE; % the index is permanent % past i [h] := \Gamma1; 8h present i [h] := \Gamma1; present i [j] := m:EQ[j]; else if begin if present i [j] ! m:EQ[j] then present i [j] := m:EQ[j]; 8h if past i [h] ! m:EQ[h] then past i [h] := \Gamma1; process the message m; when sends data to if provisional i past begin assign the index ! sn to the last checkpoint C; provisional i := FALSE; % the index is permanent % past i [h] := \Gamma1; 8h present i [h] := \Gamma1; 8h EQ i [h] := 0; packet the message % send (m) to P after first send i := TRUE; when a basic checkpoint is scheduled from if skip i then skip i := FALSE % skip the basic checkpoint else begin if provisional i then % two successive provisional indices % past begin past i [h] := \Gamma1; assign the index ! sn to the last checkpoint C; % the index is permanent % else 8h past i [h] := present i [h]; take a checkpoint C; % taking a basic checkpoint en assign the index ! sn to the last checkpoint C; provisional i := TRUE; % the index is provisional % present i [h] := \Gamma1; after first send i := FALSE; 4.5. Correctness Proof Let us first introduce the following simple observations that derive directly from the algorithm: Observation 1 For any checkpoint C i;sn;0 , there not exists a message m with m:sn - sn such that observation derives from rule take-forced(BQF) when considering C i;sn;0 is the first checkpoint with sequence number sn). Observation 2 For any checkpoint C i;sn;en , there not exists a message m with m:sn ? sn such that m is received in I i;sn;en (this observation derives from rule take-forced(BQF)). Observation 3 For any message m sent by derives from the rule send-message(BQF)). Lemma 4.1 The set with sn - 0 is a recovery line. If process P i does not have a checkpoint with index ! sn; 0 ?, the first check-point must be included in the set S. Proof If sn = 0, S is a recovery line by definition. Otherwise suppose, by the way of contradiction that S is not a recovery line. Then, there exists a message m, sent by some process P j to a process P k , that is orphan with respect to the pair (C j;sn;0 ; C k;sn;0 ). Hence, we have: C contradicts observation 1. Suppose process P k does not have a checkpoint with sequence number sn, in this case, from lemma's assumption, we replace C k;sn;0 with C k;sn 0 ;0 where sn 0 ? sn. As m is orphan wrt the pair (C received by in a checkpoint interval I k;sn 00 ;en such that m:sn ? sn 00 contradicting observation 2. Suppose process P j does not have a checkpoint with sequence number sn, in this case, from lemma's assumption, we replace C j;sn;0 with C j;sn 0 ;0 where sn 0 ? sn. As m is orphan wrt the pair (C j;sn 0 This contradicts observation 1. Hence, in all cases the assumption is contradicted and the claim follows. 2 Lemma 4.2 Let C i;sn;en i and next(C i;sn;en i ) be two local checkpoints such that C i;sn;en i Lsn If the set with en i - 0, is a recovery line L sn then the set S is a recovery line. Proof If C i;sn;en i Lsn then from definition 4.1, for each message m sent by P j such that thus no orphan message can ever exist with respect to any pair of checkpoints in S 0 . 2 From Lemma 4.1, it trivially follows: Lemma 4.3 The set L and 8j EQ i recovery line. From Lemma 4.1 and Lemma 4.2, we have each check-point belongs to at least one recovery line. In particular, belongs to all recovery lines having sequence number sn 00 such that Lemma 4.4 The set is a recovery line. Proof (Sketch) Let us assume, by the way of contradiction, S is not a recovery line. If 8j EQ i the assumption is contradicted. Otherwise, there exists a message m, sent by some process P j to a process P k , that is orphan with respect to the pair (C j;sn;EQ i and there exists a causal message chain - that brings this information to P i encoded in EQ i . Hence, we have: upon the arrival of message m, it falls in case 2 of Section 4.3. In this case, the index associated to C k;sn;EQk [j] is provisional (see the third point of Section 4.3). Before P k sends the first message m 0 forming the causal message chain -, the index has to be permanent. Hence, according to the algorithm, the index is replaced by ! sn is reset and piggybacked on m 0 . As soon as the last message of the causal message chain m 00 arrives at which is consistent by lemma 4.3. So the initial assumption is contradicted and the claim follows. 2 5. A Performance Study The Simulation Model. The simulation compares BCS, MS and the proposed algorithm (BQF) in an uniform point- to-point environment in which each process can send a message to any other and the destination of each message is a uniformly distributed random variable. We assume a system with processes, each process executes internal, send and receive operations with probability respectively. The time to execute an operation in a process and the message propagation time are exponentially distributed with mean value equal to 1 and 10 time units respectively. We also consider a bursted point-to-point environment in which a process with probability enters a burst state and then executes only internal and send events (with probability interval (when we have the uniform point-to- point environment described above). Basic checkpoints are taken periodically. Let bcf (basic checkpoint frequency) be the percentage of the ratio t=T where t is the time elapsed between two successive periodic checkpoints and T is the total execution time. For example, bcf= 100% means that only the initial local checkpoint is a basic one, while bcf= 0.1% means that each process takes 1000 basic checkpoints. We also consider a degree of heterogeneity among processes H . For example, means all processes have the same checkpoint period means 25% (resp. 75%) of processes have the checkpoint period while the remaining 75% (resp. 25%) has a checkpoint period A first series of simulation experiments were conducted by varying bcf from 0:1% to 100% and we measured the ratio Tot between the total number of checkpoints taken by an algorithm and the total number of checkpoints taken by BCS. In a second series of experiments we varied the degree of heterogeneity H of the processes and then we measured the ratio E between the total number of checkpoints taken by BQF and MS. As we are interested only in counting how many local states are recorded as checkpoints, the overhead due to the taking of checkpoints is not considered. Each simulation run contains 8000 message deliveries and for each value of bcf and H , we did several simulation runs with different seeds and the result were within 4% of each other, thus, variance is not reported in the plots. Results of the Experiments. Figure 4 shows the ratio Tot of MS and BQF in an uniform point-to-point environment. For small values of bcf (below 1.0%), there are only few send and receive events in each checkpoint interval, leading to high probability of equivalence between checkpoints. Thus BQF saves from 2% to 10% of checkpoints compared to MS. As the value of bcf is higher than 1.0%, MS and BQF takes the same number of checkpoints as the probability that two checkpoints are equivalent tends to zero. The reduction of the total number of checkpoints is amplified by the bursted environment (Figure 5) in which the equivalences between checkpoints on processes running in the burst mode are disseminated to the other processes causing other equivalences. In this case, for all values of bcf, BQF saves from a 7% to 18% checkpoints compared to MS. Performance of BQF are particularly good in an heterogeneous environment in which there are some processes with a shorter checkpointing period. These processes would push higher the sequence number leading to a very high checkpointing overhead using either MS or BCS. In Figure 6, the ratio E as a function of the degree of heterogeneity H of the system is shown in the case of uniform bursted point-to-point environment The best performance (about 30% less checkpointing than MS) are obtained when only one process has a checkpoint frequency ten times greater than the others) and 2. In Figure 7 we show the ratio Tot as a function of bcf in the case of which is the environment where BQF got the maximum gain (see Figure 6). Due to the heterogeneity, bcf is in the range between 1% and 10% of the slowest processes. We would like to remark that in the whole range the checkpointing overhead of BQF is constantly around 30% less than MS. bcf (% checkpoint period / total execution time)0.20.40.60.8 Tot total ckpt total ckpt MS Figure 4. Tot versus bcf in the uniform point-to- point environment bcf (% checkpoint period / total execution time)0.40.8 Tot total ckpt total ckpt MS Figure 5. Tot versus bcf in the bursted point-to- point environment 6. Conclusion In this paper we presented an index-based checkpointing algorithm well suited for autonomous distributed systems that reduces the checkpointing overhead compared to previous algorithms. It lies on an equivalence relation that allows to advance the recovery line without increasing its sequence number. The algorithm optimistically (and provisionally) assumes that a basic checkpoint C in a process is equivalent to the previous one in the same process by assigning a provisional index. Hence, if at the time of the first send total ckpt total ckpt MS) Figure in both the uniform point-to-point environment and the bursted point-to-point environment 1.0 3.0 5.0 7.0 9.0 bcf (% checkpoint period total execution time)0.600.80Tot total ckpt total ckpt MS Figure 7. Tot versus bcf of the slowest processes in a bursted point-to-point environment event after C that equivalence is verified, the provisional index becomes permanent. Otherwise the index is increased, as in [3, 11], and this directs forced checkpoints in other processes. We presented a simulation study which quantifies the saving of checkpoints in different environments compared to previous proposals. The price to pay is each application message piggybacks information compared to one integer used by previous algorithms. Acknowledgements . The authors would like to thank Bruno Ciciani, Michel Raynal, Jean-Michel Helary, Achour Mostefaoui and the anonymous referees for their helpful comments and suggestions. --R On Modeling Consistent Checkpoints and the Domino Effect in Distributed Systems A Communication-Induced Checkpointing Protocol that Ensures Rollback-Dependency Trackability A Distributed Domino-Effect Free Recovery Algorithm Determining Global States of Distributed Systems A Timestamp-Based Check-pointing Protocol for Long-Lived Distributed Computations A Survey of Rollback-Recovery Protocols in Message-Passing Systems Manetho: Transparent Rollback Recovery with Low Overhead Checkpointing and Rollback-Recovery for Distributed Systems Finding Consistent Global Checkpoints in a Distributed Computa- tion System Structure for Software Fault Tolerance Volatile Logging in n-Fault-Tolerant Distributed Systems Consistent Global Checkpoints that Contains a Set of Local Checkpoints --TR --CTR D. Manivannan , M. Singhal, Asynchronous recovery without using vector timestamps, Journal of Parallel and Distributed Computing, v.62 n.12, p.1695-1728, December 2002 B. Gupta , S. K. Banerjee, A Roll-Forward Recovery Scheme for Solving the Problem of Coasting Forward for Distributed Systems, ACM SIGOPS Operating Systems Review, v.35 n.3, p.55-66, July 1 2001

checkpointing;rollback-recovery;performance evaluation;global snapshot;distributed systems;timestamp management;causal dependency;fault tolerance;protocols

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Dynamically Configurable Message Flow Control for Fault-Tolerant Routing.

AbstractFault-tolerant routing protocols in modern interconnection networks rely heavily on the network flow control mechanisms used. Optimistic flow control mechanisms, such as wormhole switching (WS), realize very good performance, but are prone to deadlock in the presence of faults. Conservative flow control mechanisms, such as pipelined circuit switching (PCS), ensure the existence of a path to the destination prior to message transmission, achieving reliable transmission at the expense of performance. This paper proposes a general class of flow control mechanisms that can be dynamically configured to trade-off reliability and performance. Routing protocols can then be designed such that, in the vicinity of faults, protocols use a more conservative flow control mechanism, while the majority of messages that traverse fault-free portions of the network utilize a WS like flow control to maximize performance. We refer to such protocols as two-phase protocols. This ability provides new avenues for optimizing message passing performance in the presence of faults. A fully adaptive two-phase protocol is proposed, and compared via simulation to those based on WS and PCS. The architecture of a network router supporting configurable flow control is also described.

Introduction Modern multiprocessor interconnection networks feature the use of message pipelining coupled with virtual channels to improve network throughput and insure deadlock freedom [6,9,21,24]. Messages are broken up into small units called flits or flow control digits [9]. In wormhole switching (WS), data flits immediately follow the routing header flit(s) into the network [9]. Routing algorithms using WS can be characterized as optimistic. Network resources (e.g., buffers and channels) are committed as soon as they become available. This optimistic nature leads to high network throughput and low average message latencies. However, in the presence of I. This research was supported in part by a grant from the National Science Foundation under grant CCR-9214244 and by a grant from Spanish CICYT under grant TIC94-0510-C02-01. A preliminary version of this paper was presented in part at the 22nd Annual International Symposium on Computer Architecture, Santa Margherita Ligure, Italy, June 1995. faults, this behavior can lead to situations where the routing header can become blocked, no longer make progress, and hence cause the network to become deadlocked. Typically, additional routing restrictions and/or network resources are required to ensure deadlock freedom in the presence of faults [4,5,8,11]. For example, fault rings are constructed around convex faulty regions using additional virtual channels and attendant routing restrictions [4]. Additionally, source hardware synchronization mechanisms have been proposed to change routing decisions in the presence of faults [20], and partially adaptive routing around convex fault regions with no additional channels are feasible [5], while more recently the use of time-outs and deadlock recovery mechanisms have been proposed [22]. Alternatively, in the pipelined circuit switching (PCS) flow control mechanism, the path setup and data transmission stages are decoupled [15]. The header flit(s) is first routed to construct a path. In the presence of faults, the header may perform controlled and limited backtracking. As opposed to WS, routing algorithms based on PCS are conservative in nature, not committing data into the network until a complete path has been established. The result is an extremely robust and reliable communication protocol. However, path setup can exact significant performance penalties in the form of increased message latencies and decreased network throughput, especially for short messages. This paper proposes the use of configurable flow control mechanisms for fully adaptive routing in pipelined networks. The paper contributes dynamically configurable flow control mechanisms at the lowest level, and two-phase routing protocols at the routing layer. Routing protocols can be designed such that in the vicinity of faulty components messages use PCS style flow con- trol, where controlled misrouting and backtracking can be used to avoid faults and deadlocked configurations. At the same time messages use WS flow control in fault-free portions of the net-work with the attendant performance advantages. Such protocols will be referred to as Two-Phase protocols. A fully adaptive, deadlock-free, two-phase protocol for fault-tolerant routing in meshes and tori is proposed and analyzed in this paper. Formal properties of Two-Phase routing are established and the results of experimental evaluation are presented. The evaluation establishes the performance impact of specific design decisions, addresses the choice of conservative vs. configurable flow control for fault-tolerant routing, and discusses related deadlock/livelock freedom issues. Finally, the paper describes the architecture and operations of a single chip router for implementing Two-Phase routing protocols. The distinguishing features of this approach are, i) it does not rely on additional virtual channels over that already needed for fully adaptive routing, ii) the performance is considerably better than conservative fault-tolerant routing algorithms with equivalent reliability, iii) it is based on a more flexible fault model, i.e., supports link and/or node faults and does not require convex fault regions, iv) supports existing techniques for recovery from dynamic or transient failures of links or switches, and vi) provides routing protocols greater control over hardware message flow con- trol, opening up new avenues for optimizing message passing performance in the presence of faults. The following section introduces a few definitions, and the network, channel, and fault mod- els. A new class of flow control mechanisms is introduced in Section 3. Section 4 introduces fault tolerant routing while Section 4.1 provides an analysis of routing properties required for deadlock freedom. Section 4.2 introduces a fully adaptive two-phase routing protocol for meshes and tori. Architectural support is discussed in Section 5 and the results of simulation experiments are presented in Section 6. The paper concludes with plans for implementation of the router and future research directions. Preliminaries 2.1 Network Model Although Two-Phase routing can be used in any topology, the theoretical results are generally topology specific. The class of networks considered in this paper are the torus connected, bidirec- tional, k-ary n-cubes and multi-dimensional meshes. A k-ary n-cube is a hypercube with n dimensions and k processors in each dimension. In torus connected k-ary n-cubes, each processor is connected to its immediate neighbors modulo k in every dimension. A multidimensional mesh is similar to a k-ary n-cube, without the wrap around connections. A message is broken up into small units referred to as flow control digits or flits. A flit is the smallest unit on which flow control is performed, and represents the smallest unit of communication in a pipelined network. Each processing element (PE) in the network is connected to a routing node. The PE and its routing node can operate concurrently. We assume that one of the physical links of the routing node is used for the PE connection. The network communication links are full-duplex links, and the channel width and flit size are assumed to be equivalent. A number of virtual channels are implemented in each direction over each physical channel. Each virtual channel is realized by independently managed flit buffers, and share the physical channel bandwidth on a flit-by-flit basis. A mechanism as described in [6] is used to allocate physical channel bandwidth to virtual channels in a demand-driven manner. Flits are moved from input channel buffers to output channel buffers within a node by an internal crossbar switch. Given a header flit that is being routed through the network, at any intermediate node a routing function specifies the set of candidate output virtual channels that may be used by the message. The selection function is used to pick a channel from this set [12]. A profitable link is a link over which a message header moves closer to its destination. A backtracking protocol is one which may acquire and release virtual channels during path setup. Releasing a virtual channel that is used corresponds to freeing buffers and crossbar ports used by the message on that channel. 2.2 Virtual Channel Model The following virtual channel model is used in this paper. A unidirectional virtual channel, v i , is composed of a data channel, a corresponding channel, and a complementary channel is referred to as a virtual channel trio [15]. The routing header will traverse while the subsequent data flits will traverse . The complementary channel is reserved for use by special control flits. The corresponding channels and complementary channels essentially form a control network for coordinating fault recovery and adaptive routing of header flits including limited and controlled backtracking of header flits. The complementary channel of a trio traverses the physical channel in the direction opposite to that of its associated data channel. The channel model is illustrated in Figure 1(a). There are two virtual channels v i (v r ) and v j (v s ) from (R2) to (R1). Only one message can be in progress over a data channel. Therefore compared to existing channel models, this model requires exactly 2 extra flit buffers for each data channel - one each for the corresponding channel and complementary channel respectively. Since control flit traffic is a small percentage of the overall flit traffic, in practice all control channels across a physical link are multiplexed through a single virtual control channel [1] as shown in Figure 1(b). For example, control channel c 1 in Figure 1(b) corresponds to flit buffers v r *, v s *, v j c and v i c . 2.3 Fault Model On-line fault detection is a difficult problem. In this paper we assume the existence of fault c detection mechanisms, and focus on how such information may be used for robust, reliable com- munication. The detection mechanisms identify two different types of faults. Either the entire processing element and its associate router can fail or a communication channel may fail. When a physical link fails, all virtual channels on that particular physical link are marked as faulty. When a PE and its router fail, all physical links incident on the failed PE are also marked as being faulty. In addition to marking physical channels incident on the failed PE as being faulty, physical channels incident on PEs which are adjacent to the failed PEs and/or communication channel may be marked as unsafe. The unsafe channel [23] designation is useful because routing across them may lead to an encounter with a failed component. Some of the protocols we will present in Section 4.2 use unsafe channels. Figure 2 shows failed PEs, failed physical links and unsafe channels in a two dimensional mesh network. The failed PE can no longer send or receive any messages and thus is removed from the multi-processor network. Failures can be either static or dynamic. Static failures are present in the network when the system is powered on. Dynamic failures occur at random during operation. Both types of failures are considered to be permanent, i.e., they remain in the system until repaired. For static failures and dynamic failures that occur on idle links and routers, only header flits encounter failed links and routing protocols can attempt to find alternative paths. Figure 1. Inter-router virtual channel model a) Logical channel model for 2 virtual channels between routers R1 and b) Implementation of the logical channel model d c c d d c c d d c c d d r c d r s However, dynamic failures can occur on busy links and interrupt a message transmission. Fur- thermore, failure during the transmission of a flit across a channel can cause the flit to be lost. Since only header flits contain routing information, data flits whose progress is blocked by a failure cannot progress. They will remain in the network, holding resources, and can eventually cause deadlock. We rely on the existence of a recovery mechanism for removing such "dead" flits from the network. There exist at least two techniques for implementing distributed recovery [16, 22] under dynamic faults. In both cases, the failure of a link will generate control information that is propagated upstream and/or downstream along the message path. All resources along the path can be recovered. Alternatively, a third approach to recovering from messages interrupted by a fault can be found in [8]. All of these schemes are non-trivial, require hardware support, and have been developed elsewhere [8, 22, 16]. We will assume the existence of such a technique and evaluate its performance impact in Section 6. Scouting Switching - A family of Flow Control Mechanisms Scouting switching (SS) is a flow control mechanism that can be configured to provide specific trade-offs between fault tolerance and performance. In SS, the first data flit is constrained to remain K links behind the routing header. When the flow control is equivalent to wormhole switching, while large values can ensure path setup prior to data transmission (if a path exists). Figure 3 illustrates a time-space diagram for messages being pipelined over five links using SS mechanisms. The parameter, K, is referred to as the scouting distance or probe lead. Every time a channel is successfully reserved by the routing header, it returns a positive acknowledgment. As acknowledgments flow in the direction opposite to the Figure 2. Failed nodes and unsafe channels Faulty Node Faulty Channel Unsafe Channel routing header, the gap between the header and the first data flit can grow up to 2K - 1 links while the header is advancing. If the routing header backtracks, it must send a negative acknowledg- ment. Associated with each virtual channel is a programmable counter. A virtual channel reserved by a header increments its counter every time it receives a positive acknowledgment and it decrements its counter every time it receives a negative acknowledgment. When the value of the counter is equal to K, data flits are allowed to advance. For performance reasons, when acknowledgments are sent across the channels. In this case, data flits immediately follow the header flit. For example, in Figure 4, the header is blocked by faulty links at node A. The first data flit is constrained to remain K links behind the header at node B. From the figure, we can see that header can backtrack, releasing link A, and establish an alternate path across link C. By statically fixing the value of K, we fix the trade-off between network performance (overhead of positive and negative acks) and fault tolerance (the ability of the header to backtrack and be routed around faults). By dynamically modifying K, we can gain improved run-time trade-offs between fault tolerance and performance. If L is the message length in flits, l the number of links in the path, and K the scouting dis- tance, we can derive expressions for the minimum message latency for each type of routing mech- Figure 3. Time-space diagram of WS, Scouting, and PCS su data data data su su data Scouting Pipelined Circuit Switching Route Setup Data Transmission Routing Header PCS Acknowledgment Data Flit Scouting Acknowledgment Wormhole Switching anism. 4 Fault-tolerant Routing The basic idea proposed in this paper is for messages to be routed in one of two phases. When messages are traversing fault-free segments of the network, they are routed using protocols based on WS. When messages traverse a segment of the network with faults, a more conservative flow control mechanism, and associated fault-tolerant routing protocol is employed. The use of SS flow control to be made dynamically by simply modifying the value of K. The design of effective two-phase protocols is dependent upon the relationships between the i) scouting distance (K), ii) the number of faults (f), iii) the number of links a header flit may be forced to backtrack in routing around faults (b), and iv) the number of steps a header may be routed along non-minimal paths (m). The analysis in the following subsection establishes these relationships for k-ary n-cubes and multi-dimensional meshes. Section 4.2 describes a fully adaptive two-phase, fault-tolerant, routing protocol. 4.1 Analysis Messages are assumed to always follow shortest paths in the absence of faults. Further, when Figure 4. Backtracking out of a faulty region A Data flit progress Failed channel Routing header progress scouting l a header encounters a faulty link, it is allowed to either misroute or backtrack, with the preference given to misrouting. Theorem 1 In the absence of any previous misrouting, the maximum number of consecutive links that a header flit will backtrack over in a torus connected k-ary n-cube in a single source-destination path is is the number of faulty components. Proof: If there have been no previous misroutes, the header flit is allowed to misroute in the presence of faults even when the number of misroutes is limited. Thus, the header will only backtrack when the only healthy channel is the one previously used to reach the node (Figure 5). In the case of a k-ary n-cube, every node has 2n channels, incident on a distinct PE. Since the header arrived from a non-faulty PE, it will be forced to backtrack if 2n - 1 channels are faulty. At the next node, since the header has backtracked from a non-faulty PE and originally arrived from a non-faulty PE, it will be forced to backtrack if the remaining 2n - 2 channels are faulty. Each additional back-tracking step will be forced by 2n - 2 additional failed channels. Thus we have: Consider the second case shown in Figure 5 where there is a turn at the end of the alley. In order to cause the routing header to backtrack initially, there needs to be 2n - 1 faulty channels, the second backtrack requires 2n - 2 faulty channels while the third backtrack is necessitated by 2n - 3 node Figure 5. Node faults causing backtracking case 1 case 2 Faulty Node Faulty Link faults or 2n - 2 channel faults. All subsequent backtracks require 2n - 2 additional faults. Thus we Theorem 2 In the absence of any previous misrouting, the maximum number of consecutive links that a header flit will backtrack over in a n-dimensional mesh in a single source-destination path is is the number of faulty components. Proof: If there have been no previous misrouting operations, the message is allowed to misroute in the presence of faults, even if the maximum number of misrouting operations is limited. There are several possible cases: - The routing probe is at a node with 2n channels. This is the same case as with a torus connected k-ary n-cube. Hence, the number of faults required to force the first backtrack is 2n - 1. To force additional backtracks, 2n - 2 additional faults are required per additional backtrack. - The probe is at a node with less than 2n channels. As with the earlier cases, all channels except the one used to reach the node can be used in case of faults (either for routing or misrouting). The worst case (Figure 6(a)) occurs when the node has the minimum number of channels. In an n-dimensional mesh, nodes located at the corners only have n channels. One of the channels was used by the probe to reach the node. Hence, the failure of n - 1 channels or nodes causes the routing probe to backtrack. The probe is now on the edge of the mesh, where each node has channels. One channel was already used to reach the node the first time and another one for the previous backtracking operation, therefore, only n - 1 channels are available for routing. These channels must all be faulty to force a backtrack operation. Thus, the maximum number of mandatory backtrack operations is f div (n - 1), where f is the number of faults. - Consider the second case shown in Figure 6(b) where a turn at the end of the alley exists. In order to cause the initial backtrack, there needs to be n faults. n - 2 faults are required to cause a backtrack at the corner processing element. Each additional backtrack requires n - 1 II faults. Hence, the maximum number of backtracking operations is (f +1) div (n - 1). II. n -1 faulty channels or n - 2 faulty nodes for the first additional backtracks. The above theorems establish a relationship between the number of backtracking operations and the number of faults for both meshes and tori. Now consider the relationship between the number of misrouting operations, number of faults, and number of backtracking steps. This is determined by the configuration of faults and is specified by the following theorem. It will be useful in determining the scouting distance. Theorem 3 In a torus connected k-ary n-cube with less than 2n faults, the maximum number of consecutive backtracking steps, b, before the header can make forward progress is 3 III if Figure 6. Faults causing backtracking in a mesh Faulty Link Faulty Node (a) (b) Figure 7. Fault configuration showing required to search all inputs in one plane A Legend Source/Desination Node Failed Node Failed Channel i) the maximum number of misroutes allowed is 6, ii) misrouting is preferred over backtracking, iii) when necessary, the output channel selected by the routing function for misrouting the mes- sage, is in the same dimension as the input channel of the message. Proof: Consider Figure 7, where all of the adjacent nodes to the destination in one plane are faulty. The routing header would have to take a maximum of six misroutes to check all of the possible input links to the destination lying within a plane. This will eliminate two dimensions to search out of the n possible dimensions. If all permitted misroutes have been used or the routing header arrives at a previously visited node, the routing header must backtrack. Backtracking over a misroute removes it from the path and decrements the misroute count. The routing header backtracks two hops to point A in Figure 7. From this point, the routing header can take one misroute into any of the n - 2 remaining dimensions, j for example (where j is not one of the two dimensions forming the plane in Figure 7). The routing header is now two hops away from the node adjacent to the destination lying along dimension j. The routing header can check to see if that node is faulty with one profitable hop. If that node is faulty, then the routing header is forced to backtrack two hops back to point A. Alternatively, in two hops the header can check if the link adjacent to the destination is faulty. In this case the maximum backtrack distance is three hops back to point A. From point A, with one misroute and two profitable routes, the routing header can check the status of every node one hop away from the destination and/or every link adjacent to the destination. Since the number of faults allowed in the system is limited to 2n - 1, the existence of one healthy node and one healthy channel adjacent to the destination is guaranteed. Hence, the maximum number of backtracks that the routing header has to perform is three. Theorem 4 In a n-dimensional mesh with less than n faults, the maximum number of consecutive backtracking steps, b, before the header can make forward progress is 3 if i) the maximum number of misroutes allowed is 6, ii) misrouting is preferred over backtracking, iii) when necessary, the output channel selected by the routing function for misrouting the message is in the same dimension as the input channel of the message. Proof: Consider the case when the destination node cannot be surrounded by faults in any plane. III. If only node failures are considered, the number of backtracks required per backtracking operation is 2. Figure 8 shows the corner of a mesh where 3. At the corner node of the mesh, two of the three input/output channels of the corner node are faulty. The routing probe entering the corner node is forced to backtrack one step. However, since there cannot be any additional faulty links or nodes in the network (due to the limit in the number of faults), the routing probe can reach the destination without any further backtracking operations. If the routing probe is not at a corner node, but at a node on the edge of the mesh, then since each node on the edge of a mesh has and since a maximum of n - 1 faults are allowed, no backtracking will be required because misrouting is preferred over backtracking. Consider the case when the destination node can be surrounded by faults in some plane. This means that a situation similar to that shown in Figure 7 occurs, even in the nodes at the edge of the mesh. If the number of misroutes is limited to 6, then the results of Theorem 3 can be applied and the maximum number of consecutive backtracking steps is 3. Only 2n and n faults are required to disconnect the network in a k-ary n-cube and n-dimensional mesh respectively. However, in practice, the network can often remain connected with a considerably larger number of failed nodes and channels. If the total number of faults was allowed to be greater than 2n or n, then it is possible that some messages may be undeliverable. If allowed to remain in the network, these messages impact performance and may lead to deadlock. Techniques such as those described in 6.2Section 2.3 can be used to detect and remove such messages Figure 8. Backtracking in corner node of mesh Failed Link Failed Node from the network. 4.2 Two-Phase Routing Protocol Routing protocols operate in two phases: an optimistic phase for routing in fault-free segments and a conservative phase for routing in faulty segments. The former uses an existing fully adap- tive, minimal, routing algorithm [12]. In this section we propose two candidates for the conservative phase. The candidates differ primarily in the impact on performance as a function of the number of faults. The proposed Two-Phase (TP) protocol is shown in Figure 9 and operates as follows: In the absence of faults, TP uses a deadlock-free routing function based on Duato's Protocol (DP) [12]. In DP, the virtual channels on each physical link are partitioned into restricted and unrestricted partitions. Fully adaptive minimal routing is permitted on the unrestricted partition (adaptive while only deterministic routing is allowed on the restricted partition (deterministic channels). The selection function uses a priority scheme in selecting candidate output channels at a router node. First, the selection function examines the safe adaptive channels. If one of these channels is not available, either due to it being faulty or busy, the selection function examines the /* Structure of Two-Phase Routing */ IF detour complete THEN /* completed detour (destination reached or detour completed)*/ reset header to DP mode; END IF IF DP THEN /* route using DP routing restrictions with unsafe channels */ select safe profitable adaptive channel; RETURN; select safe deterministic channel; RETURN; IF NOT (safe deterministic channel faulty) THEN RETURN; /* blocks progress */ END IF select unsafe profitable adaptive channel; /* Acks sent or not sent depending on */ switch to SS mode & set ack counter; /* which one of the two different */ conservative phases of TP routing used */ select unsafe deterministic channel; switch to SS mode & set ack counter; set header to detour mode; END IF IF detour THEN /* route with no restrictions in detour mode */ select profitable channel; RETURN; IF #_misroutes < m THEN END IF END IF Figure 9. Structure of Two-Phase routing safe deterministic channel (if any). If the safe deterministic channel is busy, the routing header must block and wait for that channel to become free. If a safe adaptive channel becomes free before the deterministic channel is freed, then the header is free to take the adaptive channel. If the deterministic channel is faulty, the selection function will try to select any profitable adaptive channel, regardless of it being safe or unsafe. The selection function will not select an unsafe channel over an available safe channel. An unsafe channel is selected only if it is the only alternative other than misrouting or backtracking. When an unsafe profitable channel is selected as an output channel, the message enters the vicinity of a faulty network region. This is indicated by setting a status bit in the routing header. Subsequently, the counter values of every output channel traversed by the header is set to K. Values of K > 0 will permit the routing header to backtrack to avoid faults if the need arises. Message flow control is now more conservative, supporting more flexible protocols in routing around faulty regions. If no unsafe profitable channel is available, the header changes to detour mode. In detour mode, no positive acknowledgments are generated and with no positive acknowledg- ments, data flits do not advance. During the construction of the detour, the routing header performs a depth-first, backtracking search of the network using a maximum of m misroutes. Only adaptive channels are used to construct a detour. The detour is complete when all the misroutes made during the construction of the detour have been corrected or when the destination node is reached. When the detour is complete, SS acknowledgments flow again, and data flits resume progress. Note that all channels (or none) in a detour are accepted before the data flits resume progress. This is required to ensure deadlock-freedom. The detour mode is identified by setting a status bit in the header. While it is desirable to remain with WS for the fault-free routing (optimistic phase), alternatives are possible for the conservative phase. In the conservative phase of TP (Figure 9), the header enters SS mode when an unsafe channel is selected. Alternatively, in the conservative phase we may chose to continue optimistic WS flow control across unsafe channels. In this case, it not necessary to mark channels as unsafe. When WS forward progress is stopped due to faults, then detours can be constructed using increased misrouting as necessary. When a detour is completed, one acknowledgment is sent to resume the flow of the data flits. Note, in this case we always have no positive or negative acknowledgments are transmitted. When larger values of K are used (as in Figure 9), the increased ability to backtrack and route around fault regions reduces the probability of constructing detours. Thus we see that the choice of K is a trade-off between acknowledgment traffic, and the increased misrouting/backtracking that occurs in detour construction. We expect that the choice of an appropriate value of K is dependent upon the network load and failure patterns. The trade-offs are evaluated in Section 6. Note that the proofs of deadlock freedom do not rely on unsafe channels. Therefore the designer has some freedom in configuring the appropriate mechanisms as a function of the failure patterns. Figure shows a routing example using the Two-Phase routing protocol (as shown in Figure with seven node failures and initially set to 0, the routing header routes to node B where it is forced to cross an unsafe channel. The value of K is increased to 3 and the header routes profitably to node A, with the data flits advancing until node B. At node A, the routing header cannot make progress towards the destination entering detour mode, so it is misrouted upwards. After two additional misroutes can no longer be misrouted due to the limit on m. The routing header then is forced to backtrack to node A. Since there are no other output channels to select, the routing header is forced to backtrack to node C. From there, it is misrouted twice downwards and then finds profitable links to the destination. In this case, the detour is completed when the destination is reached. Also, notice that data flits do not advance while the header is in detour mode. Thus, the first data flit is still at node B. For comparison purposes, consider the use of an alternative conservative phase as described Figure 10. Routing example Failed channel Failed node above where unsafe channels are not used and K is always 0. Referring to Figure 10, the routing header is routed profitably to node A. In this case, Thus, the first data flit also reaches node A. Since it cannot be routed profitably from node A, a detour is constructed. The header is misrouted upwards three links, cannot find a path around the fault region, and therefore is forced to backtrack back to node A. The routing header is then forced to misroute to node C. From node C, it misroutes downwards, and traverses a path to the destination. Notice that in this case a path that is two hops longer since the data flits now pass through node A. However, while the header is routed from node B to node A, no acknowledgment flits are generated. These two examples indicate that the specific choice of flow control/routing protocol for the conservative phase is a trade-off that is dictated by the fault patterns and network load. The theorems in Section 4.1 cover networks with a fixed number of faults. For an arbitrary number of faults, f, small values of m, and destination node failures, it is possible that the header may backtrack to the location of the first data flit. In fact, this may occur if the links are simply busy rather than being faulty. One solution is to re-try from this point. However, it is possible that this also will not succeed. At this point, we rely on the recovery mechanism referenced in Section 2.3 to tear down the path and, if designed to do so, re-try from the source. With successive failures to establish a path from the source, some higher level protocol is relied upon to take appropriate action. This behavior particularly addresses messages destined for failed nodes. After a certain number of attempts, the higher level protocol may mark the node as unreachable from the source. While livelock is addressed in this fashion, the following theorem establishes the deadlock freedom of TP. Theorem 5 Two-Phase routing is deadlock-free. Proof: Let C be set of all virtual channels, C 1 be set of deterministic channels and C 2 be set of adaptive channels. The following situations can occur during the message routing: - If the routing header does not encounter any faulty nodes or channels, TP routing uses DP routing restrictions which have been shown to be deadlock-free in the fault-free network [12]. - If the routing header encounters an unsafe channel and selects a safe channel over the unsafe channel, then no deadlock can occur since the safe adaptive channel still is contained in the set of virtual channels C 2 and routing in this set cannot induce deadlock. - If the routing header is forced to take an unsafe adaptive channel, then no deadlock can occur since the unsafe channels are still in channel set C 2 and routing in C 2 cannot induce deadlock. - If the routing header encounters a faulty node or channel and cannot route profitably and cannot take a deterministic channel from C 1 , because it is faulty, then the routing header constructs a detour. No deadlock can occur while building the detour because the probe can always backtrack up to the node where the first data flit resides. No deadlock can occur in the attempt to construct a detour because if after several re-tries, the detour cannot be constructed, the recovery mechanism will tear down the path, thus releasing the channels being occupied by the message. - As the detour uses only adaptive channels, channels from C 2 , no deadlock can arise in routing the message after the detour has been constructed because, taking into account the condition to complete a detour, the ordering between channels in the deterministic channels, C 1 , is still preserved - Finally, the detour only uses adaptive channels from C 2 . Thus, building a detour does not prevent other messages from using deterministic channels to avoid deadlock. 5 Architectural Support Figure 11 illustrates the block diagram of a router that implements Two-Phase routing. This is a modified version of a PCS router described in [1]. Each input and output physical channel has associated with it a link control unit (LCU). The input LCU's feed a first-in-first-out (FIFO) data input buffer (DIBU) for each virtual channel. All input control channels are multiplexed over a single virtual channel and therefore feed a single FIFO control input buffer (CIBU). The data FIFO's feed the inputs of the crossbar. The control FIFO's arbitrate for access to the routing control unit (RCU). The RCU implements the two-phase routing protocol to select an output link, and maps the appropriate input link of the crossbar to the selected output link. The modified control flit is now sent out the RCU output arbitration unit to the appropriate control output virtual chan- nel. The LCUs and DIBUs support SS flow control as described later in this section. A single chip version of this router with only PCS flow control has been implemented in a metal layer CMOS process and fabricated by MOSIS [1]. The overall design contains over 14,000 transistors and is 0.311 cm square. The chip has 88 pins. The core logic of the router chip consumes 55% of the chip area and the crossbar occupies 14% of the area dedicated to the core circuitry. An additional 10% of the logic payload is devoted to the RCU. The routing header (Figure 12) for the Two-Phase protocol consists of six fields. The first field is the header bit field which identifies the flit as a routing header. The second field is the backtrack field. This bit signifies whether the routing header is going towards the source (i.e., backtracking) or towards the destination. The next field is the misroute field. It records the number of misrouting operations performed by the routing header. Since the Two-Phase protocol must be allowed a maximum of 6 misroutes to ensure the delivery of the message (in a network with up to 2n - 1 node faults), this field is three bits in size. The fourth field is the detour bit. This bit is used by the control logic to determine if the message is in detour mode. If the bit is clear and the SS bit is set, Figure 11. Overview of router chip LCU CPU CPU Data Buffer (Input/Output) Control Buffer (Input/Output) Data Input Bus Data Output Bus Control Input Bus Control Output Bus LEGEND CIBU/COBU DIBU/DOBU LCU LCU LCU LCU LCU LCU LCU CROSSBAR RCU RCU RCU INPUT Enable Buffers Figure 12. Format of header flit(s) Bit Header Back-track Misroute Detour Xn-offset X2-offset X1-offset SS the router generates an acknowledgment flit every time the routing header advances. Acknowledgments are propagated over the complementary control channel. Following the detour field is the SS bit. When the SS bit is set, SS flow control is used across every channel traversed by the header thus setting the counter to K. The next field is actually a set of offsets, one offset for each of the n dimensions in the k-ary n-cube. Their size depends on the size of the interconnection net-work (i.e., the value of k). Depending upon how the conservative phase is implemented, each physical channel will require an unsafe channel status bit maintained in the RCU. When a routing header enters the RCU, the input virtual channel address is used to access the unsafe channel store and the history store. The history store maintains a record of output channels that have been searched by a back-tracking header. Figure 13 shows the organization of the RCU. The major distinguishing features of this router architecture are due to the support for the backtracking search done by a header. A detailed discussion of architectural requirements for such routing can be found in [15, 1]. Associated with each virtual channel is a counter for recording acknowledgments and a register with the value of K, the scouting distance. For a two bit counter is required for each virtual channel. All counters are maintained in the counter management unit (CMU) in the RCU. When a positive (negative) acknowledgment flit arrives for a virtual circuit, the CMU increments (decrements) the counter that corresponds to the data virtual channel. If the counter value is K, data flits are allowed to flow. Otherwise they are blocked at the DIBU as show in Figure 14. This Figure 13. Routing control unit Channel Mappings Decision Unit Inc/Dec Banks History Store Decode Unsafe Store Header (modified) Output Virt. Chan DIBU Enable Input Virt. Chan. Header Unit Counter Management is achieved by providing DIBU output enables from the RCU. Finally, the RCU does not propagate the acknowledgment beyond the first data flit of a message. 6 Performance Evaluation The performance of the fault-tolerant protocols was evaluated with simulation studies of message passing in a 16-ary 2-cube with messages. The routing header was 1 flit long. The simulator performs a time-step simulation of network operation at the flit level. The message destination traffic was uniformly distributed. Simulation runs were made repeatedly until the 95% confidence intervals for the sample means were acceptable (less than 5% of the mean values). The simulation model was validated [14] using deterministic communication patterns. We use a congestion control mechanism (similar to [3]) by placing a limit on the size of the buffer (eight buffers per injection channel) on the injection channels. If the input buffers are filled, messages cannot be injected into the network until a message in the buffer has been routed. A flit crosses a link in one cycle. The performance of TP was compared to the performance of Duato's Protocol (DP) [12]. DP is a wormhole based routing protocol which partitions the virtual channels into two sets, adaptive and escape. The adaptive channels permit fully adaptive minimal routing while the escape channels are used to implement a deadlock-free sub-network. To measure the fault tolerance of TP, it was compared with Misrouting, Backtracking with m misroutes (MB-m) [15]. MB-m is a PCS based routing protocol which allows fully adaptive routing and up to m misroutes per virtual circuit Figure 14. Data flit flow control COBU CIBU To RCU Arb From RCU DOBU DOBU DOBU To Crossbar Enable Lines From RCU DIBU DIBU DIBU From Crossbar Router A Router B The metrics used to measure the performance of TP are average message latency and network throughput. Average message latency is the average of the time that messages spend in the net-work after their respective routing headers have been injected into the network until the time when the tail flit is consumed by the destination node. Network throughput is defined as the total number of flits delivered divided by the number of nodes in the network and the total simulation time in clock cycles. When no faults are present in the network, TP routing uses the DP routing restrictions and This results in performance that is identical to that of DP. The fault performance of TP is evaluated with a configuration of TP which uses the faulty regions, i.e., does not use unsafe channels, and then uses misrouting backtracking search to construct detours when the header cannot advance. 6.1 Static Faults Figure 15 is a plot of the latency-throughput curves of TP and MB-m with 1, 10, and 20 failed nodes randomly placed throughout the network. While the theorems developed in this paper depend on the number of faults being less than the degree of processing elements (i.e., connected k-ary n-cube), the plots show the performance of TP for larger values of faults because the faults are randomly distributed throughout the network. When randomly placed, 2n - 1 faults do not perturb the system significantly. The performance of both routing protocols drop as the number of failed nodes increase, since the number of undeliverable Figure 15. Latency-throughput of TP and MB-m with node faults Throughput Latency (Clock Cycles) Latency Vs. Throughput TP and MB-m in Faulty Network MB-m (1F) MB-m (10F) MB-m (20F) messages increases as the number of faults increase. However, the latency of TP routed messages for a given network load remains 30 to 40% lower than that of MB-m routed messages. MB-m degrades gracefully with steady but small drops in the network saturation traffic load (the saturation traffic is the network load above which the average message latency increases dramatically with little or no increase in network throughput) as the number of faults increases. Figure 16(a) shows that the latency of messages successfully routed via MB-m remains relatively flat regardless of the number of faults in the system. The number in parenthesis indicates the number of messages offered/node/5000 clock cycles. However, with the network offered load at 0.2 flits/node/cycle (30 msgs/node/5000 cycles), the latency increased considerably as the number of faults increased. This is because with a low number of faults in the system, an offered load of flits/node/cycle is at the saturation point of the network. With the congestion control mechanism provided in the simulator, any additional offered load is not accepted. However, at the saturation point, any increases in the number of faults will cause the aggregate bandwidth of the network to increase beyond saturation and therefore cause the message latency to increase and the network throughput to drop. When the offered load was at 0.32 flits/node/cycle, the network was already beyond saturation so the increase in the number of faults had a lesser effect. At low to moderate loads and with a lower number of faults, the latency and throughput characteristics of TP are significantly superior to that of MB-m. The majority of the benefit is derived from messages in fault-free segments of the network transmitting with trol). TP however, performed poorly as the number of faults increased. While saturation traffic Figure 16. Latency and throughput of TP and MB-m as function of node faults Node Failures200.0600.0Latency (Clock Cycles) Latency Vs. Node Faults TP and MB-m TP (1) MB-m (1) MB-m (30) Node Failures0.100.30Throughput Throughput Vs. Node Faults TP and MB-m TP (1) MB-m (1) MB-m (30) with one failed node was 0.32 flits/node/cycle, it dropped to slightly over with 20 failed nodes (only ~17% of original network throughput). In the simulated system (a 16- ary 2-cube), 2n - 1 faults is 3. Hence 20 failed nodes is much greater than the limit set by the theorems proposed in this paper. Figure 16 also shows the latency and throughput of TP as a function of node failures under varying offered loads. At higher loads and increased number of faults, the effect of the positive acknowledgments due to the detour construction becomes magnified and performance begins to drop. This is due to the increased number of searches that the routing header has to perform before a path is successfully established and the corresponding increase in the distance from the source node to the destination. The trade-off in this version of TP is the increased number of detours constructed vs. the performance of messages in fault-free sections of the network. With larger numbers of faults, the former eventually dominates. In this region purely conservative protocols appear to remain superior. In summary, at lower fault rates and below network saturation loads, TP performs better than the conservative counterpart. We also note that TP protocol used in the experiments was designed for 3 faults (a 2 dimensional network). A relatively more conservative version could have been configured. Figure 17 compares the performance of TP with only one fault in the network and low network traffic, both versions realize similar performance. However, with high network traffic and larger number of faults, the aggressive TP performs considerably better. This is due to the fact that with K > 0, substantial acknowledgment flit traffic can be introduced into the Throughput (Flits/Cycle/Node)100.0200.0Latency (Clock Cycles) Latency Vs. Throughput Conservative vs. Aggressive SR Aggressive (1F) Aggressive (10F) Aggressive (20F) Conservative (1F) Conservative (10F) Conservative (20F) Figure 17. Comparison of aggressive conservative SS routing behavior network, dominating the effect of an increased number of detours. 6.2 Dynamic Faults When dynamic faults occur, messages may become interrupted. In [16], a special type of control flit called, kill flit, was introduced to permit distributed recovery. When a message pipeline is interrupted, PEs that span the failed channel or PE release kill flits on all virtual circuits that were affected. These kill flits follow the virtual circuits back to the source and the destination of the messages. These control flits release any reserved buffers and notify the source that the message was not delivered, and notify the destination to ignore the message currently being received. If we are also interested in guaranteeing message delivery in the presence of dynamic faults, the complete path must be held until the last flit is delivered to the destination. A message acknowledgment sent from the destination traverses the complementary control channel, removes the path, and flushes the copy of the message at the source. Kill flits require one additional buffer in each control channel. This recovery approach is described in [16]. Here we are only interested in the impact on the performance of TP. Figure 18 illustrates the overhead of this recovery and reliable message delivery mechanism. The additional message acknowledgment introduces additional control flit traffic into the sys- tem. Message acknowledgments tend to have a throttling effect on injection of new messages. As a result, TP routing using the mechanism saturates at lower network loads and delivered messages have higher latencies. We compare the cases of i) probabilistically inserting f faults dynamically, with ii) f/2 static faults - this is the average number of dynamic faults that would occur. From the simulation results shown in Figure 18, we see that at low loads the performance impact of support for dynamic fault recovery is not very significant. However, as injection rates increase, the additional traffic generated by the recovery mechanism and the use of message acknowledgments begins to produce a substantial impact on performance. The point of interest here is that dynamic fault recovery has a useful range of feasible operating loads for TP protocols. In fact, this range extends almost to saturation traffic. 6.3 Trace Driven Simulation The true measure of the performance of an interconnection network is how well it performs under real communication patterns generated by actual applications. The network is considered to have failed if the program is prevented from completing due to undeliverable messages. Communication traces derived from several different application programs: EP (Gaussian Deviates), MM (Matrix Multiply), and MMP (another Matrix Multiply). These program traces were generated using the SPASM execution driven simulator [25]. Communication trace driven simulations were performed allowing only randomly placed physical link failures. Node failures would require the remapping of the processes, with the resulting remapping affecting performance. No recovery mechanisms were used for recovery of undeliverable messages. The traces were generated from applications executing on a 16-ary 2-cube. The simulated network was a 16-ary 2-cube with 8 and 16 virtual channels per physical link. The aggressive version of TP was used, i.e., no unsafe channels were used. Figure 19 shows three plots of the probability of completion rates for the three different program traces with differing values of misrouting (m). A trace is said to have completed when all trace messages have been delivered, hence the probability of completion is defined as the ratio of the number of traces that were able to execute to completion over the total number of traces run. If even one message cannot be deliv- ered, program execution cannot complete. The results show the effect of not having recovery mechanisms. These simulations were implemented with no re-tries attempted when a message backtracks to the source or the node containing the first data flit. This is responsible for probabilities of completion below 1.0 for even a small number of faults. The performance effect of the recovery mechanism was illustrated in Figure 18. We expect that 2 or 3 re-tries will be sufficient in practice to maintain completion probabilities of 1.0 for a larger number of faults. In some instances, an increased number of misroutes resulted in poorer completion rates. We Figure 18. Comparison of TP with and without tail-acknowledgment flits Throughput (Flits/Cycle/Node)100.0200.0Latency (Clock Cycles) Latency Vs. Throughput Comparison of Dynamic Fault-Tolerant Mechanism w/o TAck (1F) w/o TAck (10F) w/o TAck (20F) with TAck (1F) with TAck (10F) with TAck (20F) believe that this is primarily due to the lack of recovery mechanisms and re-tries. Increased misrouting causes more network resources to be reserved by a message. This may in turn increase the probability that other messages will be forced to backtrack due to busy resources. Without re- tries, completion rates suffer. We again see the importance of implementating relatively simple heuristics such as a small number of re-tries. Finally, the larger number of virtual channels offered better performance since it provided an increase of network resources and hence reduced the probability of backtracking due to busy links. 6.4 Summary of Performance Specifically, the performance evaluation provided the following insights. Link Failures0.801.00 Probability of Completion Probability of Completion M=3, V=8 M=4, V=8 M=5, V=8 Link Failures0.40.8Probability of Completion Probability of Completion M=3, V=8 M=4, V=8 M=5, V=8 Link Failures0.700.901.10 Probability of Completion Probability of Completion M=3, V=8 M=4, V=8 M=5, V=8 Figure 19. Probability of completions for various program traces and numbers of allowed misroutes . The cost of positive acknowledgments dominates the cost of detour construction, suggesting the use of low values of K, preferably . Configurable flow control enables substantial performance improvement over PCS for low to modest number of faults since the majority of traffic is in the fault-free portions, realizing close to WS performance. . For low to modest number of faults, the performance cost of recovery mechanisms is relatively low. . At very high fault rates, we still must use more conservative protocols to ensure reliable message delivery and application program completion. Conclusions Routing in the presence of faults demands a greater level of flexibility than required in fault-free networks. However, designing routers based on the relatively rare occurrence of faults, requires that all message traffic be penalized: even the messages that route through the fault-free portions of the network. Overhead may arise due to the setting up of a fault-free path prior to data transmission (PCS), marking processors, and channels faulty to construct convex fault regions [4,5], or increasing the number of virtual channels for routing messages around the faulty components [4]. From low to moderate number of faults, configurable flow control mechanisms can lead to deadlock-free fault-tolerant routing protocols whose performance is superior to more conservative routing protocols with comparable reliability. In a network with a large number of faults, TP's partially optimistic behavior results in a severe performance degradation. With conservative routing protocols, no network resources are reserved until a path has been setup between the source and the destination. TP does not require any complex renumbering scheme to provide fault-tolerance [19,20], does not require the construction of convex regions [4,5], does not require additional virtual channels [4], and the dynamic fault-tolerant version of TP does not rely on time-outs [11] or padding of messages [22]. It does, however, result in a more complex channel model which can affect link speeds. The router designed to support TP requires only slightly more hardware than a router supporting PCS [1], making the implementation very feasible. Current efforts are redesigning the PCS router for support of TP protocols. It is however apparent that one of the most important performance issues is a more efficient mechanism for implementing the positive/negative acknowledg- ments. We are currently evaluating an implementation that adds a few control signals to the physical channel, modifying the physical flow control accordingly (the logical behavior remains unchanged). By implementing acknowledgment flits in hardware, we hope to extend the superior low load performance of TP to significantly higher number of faults. --R DISHA: An efficient fully adaptive deadlock recovery scheme. A comparison of adaptive wormhole routing algorithms. The reliable router: A reliable and high-performance communication substrate for parallel computers High performance bidirectional signalling in VLSI systems. A theory of fault-tolerant routing in wormhole networks A new theory of deadlock-free adaptive routing in wormhole networks Scouting: Fully adaptive Computer Systems Performance Evaluation. The effects of faults in multiprocessor net- works: A trace-driven study Adaptive routing protocols for hypercube interconnection networks. The turn model for adaptive routing. Cray T3D: A new dimensions for cray research. Compressionless routing: A framework for fault-tolerant routing A fault-tolerant communication scheme for hypercube computers Machine abstractions and locality issues in studying parallel systems. --TR --CTR Dong Xiang, Fault-tolerant routing in hypercubes using partial path set-up, Future Generation Computer Systems, v.22 n.7, p.812-819, August 2006 Dong Xiang , Ai Chen , Jiaguang Sun, Fault-tolerant routing and multicasting in hypercubes using a partial path set-up, Parallel Computing, v.31 n.3+4, p.389-411, March/April 2005

multiphase routing;multicomputer;fault-tolerant routing;pipelined interconnection network;message flow control;virtual channels;wormhole switching;routing protocol

299336

Vexillary Elements in the Hyperoctahedral Group.

In analogy with the symmetric group, we define the vexillary elements in the hyperoctahedral group to be those for which the Stanley function is a single Schur Q-function. We show that the vexillary elements can be again determined by pattern avoidance conditions. These results can be extended to include the root systems of types A, B, C, and D. Finally, we give an algorithm for multiplication of Schur Q -functions with a superfied Schur function and a method for determining the shape of a vexillary signed permutation using de taquin.

Introduction The vexillary permutations in the symmetric group have interesting connections with the number of reduced words, the Littlewood-Richardson rule, Stanley symmetric func- tions, Schubert polynomials and the Schubert calculus. Lascoux and Sch-utzenberger [14] have shown that vexillary permutations are characterized by the property that they avoid any subsequence of length 4 with the same relative order as 2143. Macdonald has given a good overview of vexillary permutations in [16]. In this paper we propose a definition for vexillary elements in the hyperoctahedral group. We show that the vexillary elements can again be determined by pattern avoidance conditions. We will begin by reviewing the history of the Stanley symmetric functions and establishing our notation. We have included several propositions from the literature that we will use in the proof of the main theorem. In Section 2 we will define the vexillary elements in the symmetric group and the hyperoctahedral group. We state and prove that the vexillary elements are precisely those elements which avoid different patterns of lengths 3 and 4. Due to the quantity of cases that need to be analyzed we have used a computer to verify a key lemma in the proof of the main theorem. The definition of vexillary can be extended to cover the root systems of type A, B, C, and D; in all four cases the definition is equivalent to avoiding certain patterns. In Sections 3, we give an algorithm for multiplication of Schur Q-functions with a superfied Schur function. In Section 4 we outline a method for determining the shape of a signed permutation using de taquin. We conclude with several open problems related to vexillary elements in the hyperoctahedral group. Let S n be the symmetric group whose elements are permutations written in one-line notation as [w 1 generated by the adjacent transpositions oe i for Date: June 24, 1996. The first author is supported by the National Science Foundation and the University of California, Presidential Postdoctoral Fellowship. positions i and when acting on the right, i.e., be the hyperoctahedral group (or signed permutation group). The elements of are permutations with a sign attached to every entry. We use the compact notation where a bar is written over an element with a negative sign. For example is generated by the adjacent transpositions oe i for along with oe 0 which acts on the right by changing the sign of the first element, i.e., [w If w can be written as a product of the generators oe a 1 oe a 2 ap and p is minimal then the concatenation of the indices a 1 a is a reduced word for w, and p is the length of w, denoted l(w). Let R(w) be the set of all reduced words for w. The signed (or unsigned) permutations [w have the same set of reduced words. For our purposes it will useful to consider these signed permutations as the same in the infinite groups Let s - be the Schur function of shape - and let Q - be the Q-Schur function of shape -. See [15] for definitions of these symmetric functions. Definition 1. For w define the S n Stanley symmetric function by where A(D(a)) is the set of all weakly increasing sequences such that if i then we don't have a k\Gamma1 ? a k , (i.e. no descent in the corresponding reduced word). For define the C n Stanley symmetric function by where A(P (a)) is the set of all weakly increasing sequences such that if i then we don't have a no peak in the corresponding reduced word). In [20], Stanley showed that Gw is a symmetric function and used it to express the number of reduced words of a permutation w in terms of f - the number of standard tableaux of shape -, namely where ff - w is the coefficient of s - in Gw . Bijective proofs of (1.3) were given independently by Lascoux and Sch-utzenberger [13] and Edelman and Greene [4]. Reiner and Shimozono [18] have given a new interpretation of the coefficients ff - w in terms of D(w)-peelable tableaux. VEXILLARY ELEMENTS IN THE HYPEROCTAHEDRAL GROUP 3 Stanley also conjectured that there should be an analog of (1.3) for B n . This conjecture was proved independently by Haiman [7] and Kra'skiewicz [8] in the following form: where g - is the number of standard tableaux on the shifted shape -, and fi - w are the coefficients of Q - when Fw is expanded in terms of the Schur Q-functions. The Stanley symmetric functions can also be defined using the nilCoxeter algebra of S n and B n respectively(see [5] and [6]). The relationship between Kra'skiewicz's proof of (1.4) and B n Stanley symmetric functions are explored in [11]. See also [2, 10, 23] for other connections to Stanley symmetric functions. The functions Fw are usually referred to as the Stanley symmetric functions of type C because they are related to the root systems of type C. The Weyl group for the root systems of type B and C are isomorphic, so we can study the group B n by studying either root system. We extend the results of the main theorem to the root systems of type B and D at the end of Section 2. We will use these symmetric functions to define vexillary elements in S n and B n . The Stanley functions Fw can easily be computed using Proposition 1.1 below which is stated in terms of special elements in B n . There are two types of "transpositions" in the hyperoctahedral group. These transpositions correspond with reflections in the Weyl group of the root system B n . Let t ij be a transposition of the usual type i.e. be a transposition of two elements that also switches sign element s ii simply changes the sign of the ith element. Let - ij be a transposition of either type. A signed permutation w is said to have a descent at r if w r ? w r+1 . Proposition 1.1. [1] The Stanley symmetric functions of type C have the following recursive formulas: 0!i!r l(wtrs t ir )=l(w) Fwtrs t ir l(wtrss ir )=l(w) Fwtrss ir where r is the last descent of w, and s is the largest position such that w s ! w r . The recursion terminates when w is strictly increasing in which case is the partition obtained from arranging fjw decreasing order. For example, let is a descent and w and This implies wt rs 3] and we have Continuing to expand the right hand side we see [ - 4; - 3; 1; 2] is strictly increasing so Note that l(wt rs ) always equal Proposition 1.1 because of the choice for r and s. If l(wt rs - ir rs - ir rs 1. The reflections which increase the length of wt rs by exactly 1 are characterized by the following two propositions. Proposition 1.2 ([17]). If w 2 S1 or B1 and only if and no k exists such that Proposition 1.3. [1] If w 2 B1 , and i - j, then l(ws ij only if and no k exists such that either of the following are true: 2. Main Results In this section we give the definition of the vexillary elements in S n and B n . Then we present the main theorem. The proof follows after several lemmas. Definition 2. If w 2 S n then w is vexillary if Similarly, if w then w is vexillary if parts. It follows from the definition of the Stanley symmetric functions (1.3) that if w is vexillary then the number of reduced words for w is the number of standard tableaux of a single shape (unshifted for w 2 S n or shifted for For S n , this definition is equivalent to the original definition of vexillary given by Lascoux and Sch-utzenberger in [14]. They showed that vexillary permutations w are characterized by the condition that no subsequence a exists such that . This property is usually referred to as 2143-avoiding. Lascoux and Sch-utzenberger also showed that the Schubert polynomial of type A n indexed by w is a flagged Schur function if and only if w is a vexillary permutation. One might ask if the Schubert polynomials of type B, C or D indexed by a vexillary element could be written in terms of a "flagged Schur Q-function." Many other properties of permutations can be given in terms of pattern avoidance. For example, the reduced words of 321-avoiding [3] permutations all have the same content, and a Schubert variety in SL n =B is smooth if and only if it is indexed by a permutation which avoids the patterns 3412 and 4231 [9]. Also, Julian West [24] and Simion and Schmidt [19] have studied pattern avoidance more generally and given formulas for computing the number of permutations which avoid combinations of patterns. Recently, Stembridge [23] has described several properties of signed permutations in terms of pattern avoidance as well. We will define pattern avoidance in terms of the following function which flattens any subsequence into a signed permutation. VEXILLARY ELEMENTS IN THE HYPEROCTAHEDRAL GROUP 5 Definition 3. Given any sequence a 1 a of distinct non-zero real numbers, define to be the unique element such that ffl both a j and b j have the same sign. ffl for all i; j, we have jb For example, fl( - 6; 3; - 7; containing the subsequence does not avoid the pattern - 41. Theorem 1. An element w 2 B1 is vexillary if and only if every subsequence of length 4 in w flattens to a vexillary element in B 4 . In particular, w is vexillary if and only if it avoids the following patterns: This list of patterns was conjectured in [10]. Due to the large number of non-vexillary patterns in (2.1) we have chosen to prove the theorem in two steps. First, we have verified that the theorem holds for B 6 , see Lemma 2.1. Second, we show that any counter example in B1 would imply a counter example in B 6 . Lemma 2.1. Let w vexillary if and only if it does not contain any subsequence of length 3 or 4 which flattens to a pattern in (2:1). See the appendix for an outline of the code used to verify Lemma 2.1. Lemma 2.2. Let w be any signed permutation. Suppose w is a subsequence of w and let ). Then the following statements hold: 1. If the last decent of w appears in position i r 2 fi then the last descent of u will be in position r. 2. If in addition, w i s and i s is the largest index in w such that this is true then s is the largest index in u such that this is true. 3. If then and - jk are transpositions of the same type. One can check the facts above follow directly from the definition of the flatten function. Lemma 2.3. For any v 2 B1 and any 0 there exists an index k such that Similarly, if there exists an index k such that either 6 SARA BILLEY AND TAO KAI LAM Proof. If l(vt ir pick k such that v k is the largest value in fv k rg. Then no j exists such that k Say l(vs ir r such that v k chose k such that v k is the largest value in fv k rg. Then no j exists such that k 1. On the other hand, if no such k exists, then choose k such that v k is the smallest value in fv k ? \Gammav ig. Then no exists such exists such that \Gammav r Lemma 2.4. Given any w 2 B1 and any subsequence of w, say w Similarly, if Proof. If so since the flatten map preserves the relative order of the elements in the subsequence and signs. Therefore, l(vt jk If exists such that w i j . This in turn implies that no exists such that and \Gammaw i k so since the flatten map preserves the relative order of the elements in the subsequence and signs. Also, if Therefore, l(vs jk exists such that , and no exists such that \Gammaw i k . This in turn implies that no exists such that \Gammav exists such that Lemma 2.5. Given any w 2 B1 , if w is non-vexillary then w contains a subsequence of length 4 which flattens to a non-vexillary element in B 4 . Proof. Since w is non-vexillary then either Fw expands into multiple terms on the first step of the recurrence in (1.5) or else non-vexillary. Assume the first step of the recurrence gives other terms 1g be an order preserving map onto the 4 smallest distinct numbers in the range. Let w fore, the recursion implies Hence, w is not vexillary, and it follows that w contains the non-vexillary subsequence If on the other hand the first step of the recursion gives rs - ir and v is not vexillary. Assume, by induction on the number of steps until the recurrence branches into multiple terms that v contains a non-vexillary subsequence say v a v b v c v d . If VEXILLARY ELEMENTS IN THE HYPEROCTAHEDRAL GROUP 7 is exactly the same non-vexillary subsequence. So we can assume the order of the set fa; b; c; d; than or equal to 6. Let be an order preserving map which sends the numbers 1 through 6 to the 6 smallest distinct integers in the range. Let w contains a non-vexillary subsequence, hence v 0 is not vexillary by Lemma 2.1. We will use the recursion on Fw 0 to show that w 0 is not vexillary in B 6 . From Lemma 2.2 it follows that By Lemma 2.3, rs possibly other terms. Irregardless of whether there are any other terms in expansion of Fw 0 , w 0 is not vexillary since v 0 is not vexillary. Again by Lemma 2.1, this implies w 0 contains a non-vexillary subsequence of length 4, say w h . Hence, w contains the non-vexillary subsequence This proves one direction of Theorem 1. Lemma 2.6. Given any w 2 B1 , if w contains a subsequence of length 4 which flattens to a non-vexillary element in B 4 then w is non-vexillary. Proof. Assume w is vexillary then let w be the sequence of signed permutations which arise in expanding using the recurrence (1.5). This recurrence terminates when the signed permutation w (k) is strictly increasing, hence w (k) does not contain any of the patterns in (2.1). Replace w by the first w (i) such that w (i) contains a non-vexillary subsequence and w (i+1) does not, and let Say w a w b w c w d is a non-vexillary subsequence in w. If would be exactly the same non-vexillary subsequence. This contradicts our choice of v. So we can assume that the order of the set fa; b; c; d; than or equal to 6. As in the proof of Lemma 2.5, let be an order preserving map onto the smallest 6 distinct numbers in the range. Let To simplify notation, we also contains a non-vexillary subsequence hence w 0 is not vexillary by Lemma 2.1. As in 2.5 one can show contains a non-vexillary subsequence and v 0 does not there must be another term in Fw 0 indexed by a reflection - should note that it is possible that i must be different types of transpositions. 1.3 and the definition of the flatten function, we have l(wt rs - jr rs ) ? 0. By Lemma 2.3 there exists a reflection - kr such that l(wt rs - kr rs We must have - kr 6= - ir since - possibly other terms. This proves w is not vexillary contrary to our assumption. This completes the proof of Theorem 1. The definition of vexillary can be extended to Stanley symmetric functions of type B and D. These cover the remaining infinite families of root systems. For these cases, we define vexillary to be the condition that the function is exactly one Schur P -function. The signed permutations which are B and D vexillary can again be determined by avoiding certain patterns of length 4. Theorem 2. An element w 2 B1 is vexillary for type B if and only if every subsequence of length 4 in w flattens to a vexillary element of type B in B 4 . In particular, w is vexillary if and only if it avoids the following patterns: An element w 2 D1 is vexillary for type D if and only if every subsequence of length 4 avoids the following patterns: Note, that the patterns that are avoided by vexillary elements of type D are not all type D signed permutations but instead include some elements with an odd number of negative signs. The proof of Theorem 2 is very similar to the proof of Theorem 1 given above. We omit the details in this abstract. 3. A rule for multiplication Lascoux and Sch-utzenberger noticed that the transition equation for Schubert polynomials of vexillary permutations can be used to multiply Schur functions [16][p.62]. L. Manivel asked if the transition equations for Schubert polynomials of types B, C, and could lead to a rule for multiplying Schur Q-functions. The answer is "sometimes". There are only certain shifted shapes - which can easily be multiplied by an arbitrary Schur Q-function. Therefore, we have investigated a different problem. In this section VEXILLARY ELEMENTS IN THE HYPEROCTAHEDRAL GROUP 9 we present an algorithm for multiplication of a Schur Q-function by a superfied Schur Let OE be the homomorphism from the ring of symmetric functions onto the subring generated by odd power sums defined by The image of a Schur function under this map, OE(s - ), is called a superfied Schur function. The superfied Schur functions appear in connection with the Lie super algebras [21][25]. The Stanley symmetric functions of type A and C which are indexed by permutations are related via the superfication operator. Proposition 3.1. [2][11][22] For v 2 S n , we have F be any signed permutation. We denote the signed permutation v. Also, if signed permutation, let w \Theta v be [w Lemma 3.2. For v 2 S1 and w 2 B1 we have Proof. From (1.2), when v 2 S1 , F v is equal to F 1 n \Thetav since a 1 a only assuming n is large enough, the reduced words for w \Theta v are all shuffles of a reduced word for w with a reduced word for v. One can check that the admissible monomials in F v\Thetaw are exactly the product of admissible monomials in Fw and F 1 n \Thetav counted with their coefficients. From Lemma 3.2 and Proposition 3.1 we have the following corollary. Corollary 3.3. Let w 2 B1 such that Then and Fw\Thetav can be determined by the recursive formula in Proposition 1.1. We remark that Corollary 3.3 can be used to multiply two Schur Q-functions in the special case that one of the shapes is equivalent to a rectangle under jeu de taquin. In this case, if a shifted shape - is equivalent to a rectangle ae then Q [12]. For each straight shape ae one can easily choose a permutation v in S1 with that shape. However, for any straight shape - other than a rectangle, the expansion of OE(s - ) will always be a sum of more than one Q - . Since F only if v 2 S1 , the algorithm for multiplying FwF v given above will not carry over for arbitrary elements of B1 . 4. The shape of a signed permutation Given a vexillary element w, for which straight shape - does for which shifted shape - does there are several ways to determine this shape: the transition equation [16, p. 52], inserting a single reduced word using the Edelman-Greene correspondence [4], or by rearranging the code in decreasing order [20]. Similarly, for vexillary elements of type C one can find this shape for the signed permutation by using the recursive formula (1.5) or by using the Kra'skiewicz insertion [8] or Haiman procedures [7] on a single reduced word. There is another method for computing the shape of a C n -vexillary element using jeu de taquin. We describe this method below. For any standard Young tableau U of shifted shape and any standard Young tableau V of straight shape we form a new standard shifted tableau U V by jeu de taquin as follows: 1. Embed U into the shifted shape 2. Obtain a tableau R by filling the remaining boxes of ffi with starting from the rightmost column and in each column from bottom to top. 3. Add j-j to each entry of V to obtain S. 4. Append R on the left side of S to obtain T . 5. Delete the box containing 1 0 in T . If the resulting tableau is not shifted, apply jeu de taquin to fill in the box. Repeat the procedure for the box containing 2 0 and so on until all the primed numbers are removed. 6. The resulting tableau of shifted shape is denoted U V . We illustrate the procedure with an example. Let Then, Steps 1 through 4 will produce the following tableau: Deleting the boxes and applying jeu de taquin as in Step 5 gives U Note that different choices for V will result in different shapes for U V . If V is chosen to be the the standard tableau with entries in the first row, in the second row etc., we will say U V has shape - -. So, the result of combining - and - by jeu de taquin in the example is the shape (4; 2) (2; There is a canonical decomposition of any signed permutation into the product of a signed permutation and an unsigned permutation with nice properties. Let w be an element of B n , not necessarily vexillary. Rearrange the numbers in w in increasing order VEXILLARY ELEMENTS IN THE HYPEROCTAHEDRAL GROUP 11 and denote this new signed permutation by u. Let v 2 S n so that uv. Note that and v is [4; 2; Definition 4. Given an element w of S n , the code of w is defined to be the composition g. The shape of w is defined to be the transpose of the partition given by rearranging the code in decreasing order. It is well known that if - is the shape of S n vexillary permutation v then G Furthermore, for each standard tableau Q of shape - there exists a unique reduced word for v with recording tableau Q under the Edelman-Greene correspondence [4]. Also, the reduced words for u are in bijective correspondence with the shifted standard tableau of shape -. Recall from Proposition 1.1 that if u is a strictly increasing signed permutation, we have F is the strictly decreasing sequence given by fju is the same set as fjw 0g. Therefore, it is easy to determine the shape of u. Definition 5. For any w 2 B n , let and u is strictly increasing. Let - be the shape of u, and let - be the shape of v as an element of . Define the shape of w , denoted -(w), to be the shape - -. In the case of vexillary signed permutations, we claim that is the shape of w. This is true in the case when w has all its numbers in increasing order. Before we prove the claim, here are some results that we will need. We refer the reader to the references for their proofs. Proposition 4.1. [2][11][22] Let w be a signed permutation with no positive numbers. -. The shape of w is Proposition 4.2. [11, Theorem 3.24] Let a 1 a 2 a am be a reduced word for some signed permutation w in B n . Suppose Q and Q 0 are the recording tableaux that are obtained by applying Kra'skiewicz insertion on a 1 a 2 a am and a 2 a also be obtained by deleting the box containing the entry 1 in Q and applying jeu de taquin to turn Q into a shifted tableau and subtracting 1 from every entry. Theorem 3. For , the shape of w, -(w) is the shape of some Kra'skiewicz recording tableau for some reduced word of w. Hence Q -(w) appears in the expansion of Fw with a non-zero coefficient. Proof. Let uv be the canonical decomposition of w. Given a reduced word a, denote the Kra'skiewicz recording tableau by K(a). Fix a reduced word b of u with recording tableau U = K(b) under Kra'skiewicz insertion. Let V be the standard tableau of straight shape with entries filled sequentially in rows from left to right, top to bottom. Let c 2 R(v) be the reduced word which inserts under the Edelman-Greene correspondence to V . We will actually show that the reduced word rise to the Kra'skiewicz recording tableau U V which by definition has the same shape as -(w). We prove the claim by induction on the number of positive numbers in w. When there is no positive number in w, this follows from Proposition 4.1 and there are no jeu de taquin slides necessary. Now suppose w has p positive numbers. Let be the signed permutation that is the same as w except that m is signed. We can write w 0 as a product is the arrangement of w 0 in increasing order and v is the same as in the decomposition for w. Let Note that u with a be a reduced word for z, then abc 2 R(w 0 ). Let U 0 be the recording tableau for ab 2 R(u 0 ). positive numbers, by the induction hypothesis, the recording tableau of the Kra'skiewicz insertion of abc 2 R(w 0 ) is given by U 0 V . From Proposition 4.2, the tableau U can be obtained from U 0 by deleting m boxes labeled applying jeu de taquin to fill them up. Since u and u 0 are strictly increasing, by Proposition 1.1, we know their shapes explicitly. Note that u 0 was chosen so that the m boxes which are vacated from the shape for u 0 in the jeu de taquin process form a vertical strip in the (n \Gamma p)th column (since w 2 B n ). Therefore, filling these m boxes with the next higher primed numbers we obtain the tableau T which appears in Step 4 when computing -(w). Therefore, since U 0 V gives K(abc) and evacuation of a from U 0 gives U then U V must give K(bc). Therefore, U V is the recording tableau for bc 2 R(w). From this proof one can also prove the following corollaries. Corollary 4.3. For any vexillary element w 2 B n , we have where -(w) is the shape of w as in Definition 5. Corollary 4.4. Given any w 2 B n , let uv be the canonical decomposition of w. For any a 2 R(u) and any b 2 R(v) with recording tableaux U and respectively, then U V is the recording tableau for the reduced word ab 2 R(w). 5. Open Problems The vexillary permutations in S n have many interesting properties. We would like to explore the possibility that these properties have analogs for the vexillary elements in 1. Is there a relationship between smooth Schubert varieties in SO(2n + 1)=B and vexillary elements? In particular, does smooth imply vexillary as in the case of S n ? 2. Is there a way to define flagged Schur Q-functions so that the Schubert polynomial indexed by w of type B or C is a flagged Schur Q-functions if and only if w is vexillary? VEXILLARY ELEMENTS IN THE HYPEROCTAHEDRAL GROUP 13 3. Are there other possible ways to define vexillary elements in B n so that any of the above questions can be answered? Below is a portion of the LISP code used to verify Theorem 1 for B 6 . The calculation was done on a Sparc 1 by running (grind-patterns 6 'c). (setf *avoid-patterns* (list (defun grind-patterns (n type) (flet ((helper (perm) (if (not (eq (and (avoid-subsequences perm (avoid-subsequences perm 4)) (vex-p perm type))) (format t "ERROR::NEW PATTERN: ~a ~%" perm) (format t ".")))) (all-perm-tester n #'helper (defun avoid-subsequences (the-list size) (let ((results t)) (catch 'foo (flet ((helper (tail) (when (member (flatten-seq tail) *avoid-patterns* :test #'equal) (setf results nil) (throw 'foo nil) (all-subsequences-tester (reverse the-list) size #'helper nil)) (throw 'foo t)) --R Transition Equations for Isotropic Flag Manifolds Schubert polynomials for the classical groups Some Combinatorial Properties of Schubert Polyno- mials Schubert Poly- nomials Schubert Polynomials and the NilCoxeter Algebra Dual equivalence with applications Criterion for smoothness of schubert varieties in SL(n) Bn Stanley Structure de hopf de l'anneau de cohomologie et de l'anneau de grothendieck d'une variete de drapeaux Oxford University Press The geometry of flag manifolds Algebraic Combin. of Combinatorics On the number of reduced decompositions of elements of Coxeter groups personal communication Permutations with forbidden sequences A theory of shifted Young tableaux --TR --CTR Bridget Eileen Tenner, On expected factors in reduced decompositions in type B, European Journal of Combinatorics, v.28 n.4, p.1144-1151, May, 2007 S. Egge , Toufik Mansour, 132-avoiding two-stack sortable permutations, Fibonacci numbers, and Pell numbers, Discrete Applied Mathematics, v.143 n.1-3, p.72-83, September 2004

stanley symmetric function;reduced word;hyperoctahedral group;vexillary