1 Introduction
Given the nowpervasive nature of parallelism in computation, there has been a tremendous amount of research into efficient parallel algorithms for a wide range of tasks. A popular approach has been to map existing sequential algorithms to parallel architectures, by exploiting their inherent parallelism. In this paper, we will focus on two specific variants of this strategy.
The deterministic approach, e.g. [5, 7, 18, 6, 25, 8] has been to study the directedacyclic graph (DAG) step dependence in classic, widelyemployed sequential algorithms, showing that, perhaps surprisingly, this dependence structure usually has low depth. One can then design schedulers which exploit this dependence structure for efficient execution on parallel architectures. As the name suggests, this approach ensures deterministic outputs (i.e. outputs uniquely determined by the input), and can yield good practical performance [7], but requires a nontrivial amount of knowledge about the problem at hand, and the use of carefullyconstructed parallel schedulers [7].
To illustrate, let us consider the classic sequential greedy strategy for solving the maximal independent set (MIS) problem on arbitrary graphs: the algorithm examines the set of vertices in the graph following a fixed, random sequential priority order, adding a vertex to the independent set if and only if no neighbor of higher priority has already been added. The basic insight for parallelization is that the outcome at each node may only depend on a small subset of other nodes, namely its neighbors which are higher priority in the random order. Blelloch, Fineman and Shun [7] performed an indepth study of the asymptotic properties of this dependence structure, proving that, for any graph, the maximal depth of a chain of dependences is in fact
with high probability, where
is the number of nodes in the graph. Recently, an impressive analytic result by Fischer and Noever [13] provided tight bounds on the maximal dependency depth for greedy sequential MIS, effectively closing this problem for MIS. Beyond greedy MIS, there has been significant progress in analyzing the dependency structure of other fundamental sequential algorithms, such as algorithms for matching [7], list contraction [25], Knuth shuffle [25][5], and graph connectivity [5].An alternative approach has been to employ relaxed data structures to schedule taskbased programs. Starting with Karp and Zhang [19], the general idea is that, in some applications, the scheduler can relax the strict order induced by following the sequential algorithm, and allow tasks to be processed speculatively ahead of their dependencies, without loss of correctness. A standard example is parallelizing Dijkstra’s singlesource shortest paths (SSSP) algorithm, e.g. [3, 22, 20]
: the scheduler can retrieve vertices in relaxed order without breaking correctness, as the distance at each vertex is guaranteed to eventually converge to the minimum. The tradeoff is between the performance gains arising from using simpler, more scalable schedulers, and the loss of determinism and the wasted work due to relaxed priority order. This approach is quite popular in practice, as several highperformance relaxed schedulers have been proposed, which can attain stateoftheart results in settings such as graph processing and machine learning
[20, 14]. At the same time, despite good empirical performance, this approach still lacks analytical bounds, and results are no longer deterministic.In this paper, we ask: is it possible to achieve both the simplicity and good performance of relaxed schedulers as well as the predictable outputs and runtime upper bounds of the “deterministic” approach?
Contribution.
In a nutshell, this work shows that a natural family of fair relaxed schedulers—providing upper bounds on the degree of relaxation, and on the number of inversions that a task can experience—can execute a range of iterative sequential algorithms deterministically, preserving the dependence structure, and provably efficiently, providing analytic upper bounds on the total work performed. Our results cover the classic greedy sequential graph algorithms for maximal independent set (MIS), matching, and coloring, but also algorithms for list contraction and generating permutations via the Knuth shuffle. We call this class iterative algorithms with explicit dependencies. Our main technical result is that, for MIS and matching in particular, the overhead of relaxed scheduling is independent of the graph size or structure. This analytical result suggests that relaxed schedulers should be a viable alternative, a finding which is also supported by our preliminary concurrent implementation.
Specifically, we consider the following framework. Given an input, e.g., a graph, the sequential algorithm defines a set of tasks, e.g. one per graph vertex, which should be processed in order, respecting some fixed, arbitrary data dependencies, which can be specified as a DAG. Tasks will be accessible via a scheduler, which is relaxed, in the sense that it could return tasks out of order. This induces a sequential model,^{1}^{1}1We consider this sequential model, similar to [7], since there currently are no precise ways to model the contention experienced by concurrent threads on the scheduler. Instead, we validate our findings via a fully concurrent implementation. where at each step, the scheduler returns a new task: for simplicity, assume for now that the scheduler returns at each step a task chosen uniformly at random among the top available tasks, in descending priority order. (We will model realistic relaxed schedulers [21, 2] precisely in the following section.)
Assume a thread receives a task from the scheduler. Crucially, the thread cannot process the task if it has data dependencies on higherpriority tasks: this way, determinism is enforced. (We call such a failed delete attempt by the thread a wasted step.) However, threads are free to process tasks which do not have such outstanding dependencies, potentially outoforder (we call these successful steps.) We measure work in terms of the total number of scheduler queries needed to process the entire input, including both successful and unsuccessful removal steps.
We provide a simple yet general approach to analyze this relaxed scheduling process, by characterizing the interaction between the dependency structure induced by a given problem on an arbitrary input, and the relaxation factor in the scheduling mechanism, which yields bounds on expected work when executing such algorithms via relaxed schedulers. Our approach extends to general iterative algorithms, as long as task dependencies are explicit, i.e., can be statically expressed given the input, and tasks can be randomly permuted initially.
The work efficiency of this framework will critically depend on the rate at which threads are able to successfully remove dependencyfree tasks. Intuitively, this rate appears to be highly dependent on (1) the problem definition, (2) the scheduler relaxation factor , but also on (3) the structure of the input. Indeed, we show that in the most general case, a relaxed scheduler can process an input described by a dependency graph on nodes and edges and incur wasted steps, i.e. total steps. This result immediately implies a low “cost of relaxation” for problems whose dependency graph is inherently sparse, such as Greedy Coloring on sparse graphs, Knuth Shuffle and List Contraction, which are characterized by a dependency structure with only edges. Hence, in general, such sparse problems incur negligible relaxation cost when .
Our main technical result is a counterintuitive bound for greedy MIS: our framework equipped with a relaxed scheduler can execute greedy MIS on any graph and experience only wasted steps (i.e. total steps), regardless of the size or structure of . This result is surprising as well as technically nontrivial, and demonstrates that for MIS on large graphs, operationlevel speedups provided by relaxation come with a negligible global tradeoff. A similar result holds for maximal matching.
In the broader context of the parallel scheduling literature, our results suggest that task priorities can be supported in a scalable manner, through relaxation, without loss of determinism or work efficiency. We believe this is the first time this observation is made. We validate our results empirically, via a preliminary implementation of the scheduling framework in C++, based on a lockfree extension of the MultiQueue relaxed schedulers [21]. Our broad finding is that this relaxed scheduling framework can ensure scalable execution, with minimal overheads due to contention and verifying task dependencies. For MIS on large graphs, we obtain a solution, with 6x speedup at threads versus an optimized sequential baseline.
Related Work.
Our work is inspired by the line of research by Blelloch et al. [5, 7, 18, 6, 25], as well as [12, 11, 13], whose broad goal has been to examine the dependency structure of a wide class of iterative algorithms, and to derive efficient scheduling mechanisms given such structure.
At the same time, there are several differences between these results and our work. First, at the conceptual level, [7, 25] start from analytical insights about the dependency structure of algorithms such as greedy MIS, and apply them to design scheduling mechanisms which can leverage this structure, which require problemspecific information. In some cases, e.g. [7], the scheduling mechanisms found to perform best in practice differ from the structure of the schedules analyzed. By contrast, we start from a realistic model of existing highperformance relaxed schedulers [21], and show that such schedulers can automatically and efficiently execute a broad set of iterative algorithms. Second, at the technical level, the methods we develop are different: for instance, the fact that the iterative algorithms we consider have low dependency depth [5, 7, 25] does not actually help our analysis, since a sequential algorithm could have low dependency depth and be inefficiently executable by a relaxed scheduler: the bad case here is when the dependency depth is low (logarithmic), but each “level” in a breadthfirst traversal of the dependency graph has high fanout. Specifically, we emphasize that the notion of prefix defined in [7] to simplify analysis is different from the set of positions which can be returned by the relaxed stochastic scheduler: for example, the parallel algorithm in [7] requires each prefix to be fully processed before being removed, whereas acts like a sliding window of positions in our case. The third difference is in terms of analytic model: references such as [7] express work bounds in the CRCW PRAM model, whereas we count work in terms of number of tasks processing attempts. Our analysis is sequential, and we implement our algorithms on a shared memory architecture to demonstrate empirically good performance.
To our knowledge, the first instance of a relaxed scheduler is in work by Karp and Zhang [19], for parallelizing backtracking strategies in a (synchronous) PRAM model. This area has recently become extremely active, with several such schedulers (also called relaxed priority queues) being proposed over the past decade, see [24, 4, 26, 3, 15, 20, 21, 2, 23] for recent examples. In particular, we note that stateoftheart packages for graph processing [20] and machine learning [14] implement such relaxed schedulers.
Recent work by a subset of the authors [2] showed that a simple and popular priority scheduling mechanism called the MultiQueue [21, 15, 14] enforces strong probabilistic guarantees on the rank of elements returned, in an idealized model. Concurrent work [1] proves that these guarantees in fact hold in asynchronous concurrent executions, under some analytic assumptions. Based on this result, our work bounds should hold when using MultiQueues as relaxed schedulers, in concurrent executions.
Parallel scheduling [10, 9] is an extremely vast area and a complete survey is beyond our scope. We do wish to emphasize that standard workstealing schedulers will not provide this type of work bounds, since they do not provide any guarantees in terms of the rank of elements removed: the rank becomes unbounded over long executions, since a single random queue is sampled at every stealing step [2]. To our knowledge, there is only one previous attempt to add priorities to workstealing schedulers [17], using a multilevel global queue of tasks, partitioned by priority. This technique is different, and provides no work guarantees.
2 Executing Iterative Algorithms via Priority Schedulers
2.1 Modeling Relaxed Priority Schedulers
In the following, we will provide the sequential specification of a generic relaxed priority scheduler , which contains a set of pairs. A relaxed priority scheduler will provide the following methods:

(), which returns a pair and deletes it from the structure, if a task is available, or , otherwise. The relaxation guarantees of this operation are precisely defined below;

(), which returns whether the scheduler still has tasks or not;

( ), which inserts a new task into .
Let be the rank of the task which is returned by the operation, among all tasks present in . We say that task experiences a priority inversion at an step if a task of lower priority than is retrieved at that step. For any task , let be the number of inversions which the task experiences before being removed.
Definition 1.
Fix a relaxed priority scheduler , with parameters , the rank bound, and , the fairness bound. We say that is an relaxed priority scheduler if it ensures the following:

Rank Bound. For any time , and any integer ,
. 
Fairness Bound. For any task , and any integer , .
Relation to Practical Schedulers.
Upon inspection, both the SprayList [3] and the MultiQueue [21] relaxed priority schedulers ensure these exponential tail bounds on both rank and fairness, under some analytic assumptions. These conditions are trivially ensured by deterministic implementations such as [26]. In particular, the SprayList ensures these bounds with parameters and in , where is the number of processors [3]. MultiQueues ensure these bounds with parameters , and , where is the number of distinct priority queues [2]. This holds even in concurrent executions [1].
In the following, it will be convenient to assume a single parameter , which upper bounds both the rank and the relaxation parameters. We call the relaxed scheduler simply a relaxed scheduler.
2.2 A General Scheduling Framework
We now present our framework for executing taskbased sequential programs, whose pseudocode is given in Algorithm 1. We assume a permutation which dictates an execution order on tasks. If is the element in , we will write and ( for label). Algorithm 1 encapsulates a large number of common iterative algorithms on graphs, including Greedy Vertex Coloring, Greedy Matching, Greedy Maximal Independent Set, Dijkstra’s SSSP algorithm, and even some algorithms which are not graphbased, such as List Contraction and Knuth Shuffle [5]. We show sample instantiations of the framework in Section 2.3.
Algorithm 2 gives a method for adapting Algorithm 1 to use a relaxed queue, given an explicit dependency graph whose nodes are the tasks, and whose edges are dependencies between tasks. Importantly, given the dependency graph , Algorithm 2 gives the same output as Algorithm 1, irrespective of the relaxation factor . As usual, we write and . We assume that the permutation represents a priority order so that an edge means that depends on if and viceversa. In the former case, we say that is a successor of and is a predecessor of .
Our main result regarding Algorithm 2, proven formally in Section 3.1, argues that if is chosen uniformly at random from among all vertex permutations, then Algorithm 2 completes in at most iterations (compared to exactly for Algorithm 1). This result demonstrates that provided is not too dense, the “cost of relaxation” is low for the class of problems which admit uniformly random task permutations. Notably, this class includes all of the problems mentioned above, except for Dijkstra’s algorithm (since there, needs to respect the ordering of nodes sorted by distance from the source).
2.3 Example Applications
Applying the sequential taskbased framework of Algorithm 1 only requires an implementation of . Implementing the relaxed framework in Algorithm 2 further requires (either explicitly or via a predecessor query method). We now give examples for Greedy Vertex Coloring and List Contraction, whose dependency graph is implicit.
Greedy Vertex Coloring.
Vertex Coloring is the problem of assigning a color (represented by a natural number) to each vertex of the input graph, , such that no adjacent vertices share a color. The Greedy Vertex Coloring algorithm simply processes the vertices in some permutation order, , and assigns each vertex in turn the smallest available color. The implementation of for Greedy Vertex Coloring needs to determine the color of , which can be done as described below:
Since the underlying dependency graph is just the input graph with edge orientations given by , this is all that needs to be provided.
List Contraction.
List Contraction takes a doubly linked list, , and iteratively contracts its nodes. Contracting a node consists of swinging two pointers: and , effectively removing from the list. List Contraction is useful, e.g., for cycle counting. Although List Contraction is not inherently a graph problem, we can still construct a dependency graph whose nodes are list elements and with an edge between elements which are adjacent in . If we induce a priority order on list elements (e.g. uniformly at random), then there is an induced orientation of the edges of the dependency graph which forms a DAG. Then a predecessor query on consists of checking whether either or is an unprocessed predecessor. can be implemented with just the two steps of contraction above (possibly along with the metrics the application is computing).
2.4 Greedy Maximal Independent Set
We give a variant of Algorithm 2 adapted for Greedy Maximal Independent Set (MIS), which makes use of some exploitable substructure. In particular, once some neighbor, , of a vertex is added to the MIS, then can never be added to the MIS, at which point ’s dependents no longer have to wait for to be processed. Algorithm 4 implements MIS in the framework of Algorithm 2 while also making use of this observation. Interestingly, Algorithm 4 can also be used to find a maximal matching by taking the input graph of the matching instance and converting it to a graph , where has a vertex for each edge in and there is an edge between vertices of if the corresponding edges of share an incident vertex. (One can view matching as an “independent set” of edges, no two of which are incident to the same vertex.)
3 Analysis
In this section, we will bound the relaxation cost for the general framework (Algorithm 2) and for Maximal Independent Set (Algorithm 4). Algorithm 2 is easier to analyze and will serve as a warmup. Note that in both cases, iterations are required to process all nodes and are necessary even with no relaxation. Thus, we can think of the “cost” of relaxation as the number of further iterations beyond the first , which can be equivalently counted as the number of reinsertions performed by the algorithm. We will sometimes refer to executing such a reinsertion as a “failed delete” by .
Our primary goal will be to bound the number of iterations of the for loops in Algorithm 2 and 4 when running them sequentially with a relaxed priority queue. Although the initial analysis is sequential, the algorithms are parallel: threads can each run their own for loops concurrently and correctness is maintained. The difficulty in extending the analysis to the asynchronous setting is that it is not clear how to model failed deletes of dependents of a node that is being processed. The likelihood of such deletes depend on particulars of both the problem (i.e. how long processing and dependency checking steps actually take) and the thread scheduler and so are hard to model in our generic framework. Instead, we show empirically that our bounds hold in practice on a realistic asynchronous machine where threads run the loops fully in parallel.
The theorems we will prove are the following. Given a dependency graph with vertices and edges, we first bound the number of iterations of Algorithm 2:
Theorem 1.
Algorithm 2 runs for iterations in expectation.
By contrast to Algorithm 2, we show that using a relaxed queue for computing Maximal Independent Sets on large graphs has essentially no cost at all, even for dense graphs! In particular, Algorithm 4 incurs a relaxation cost with no dependence at all on the size or structure of , only on the relaxation factor :
Theorem 2.
Algorithm 4 runs for iterations in expectation.
Before delving into the individual analyses, we first consider some key characteristics of a particular relaxed queue which will be at play, and quantify them in terms of the fairness and rank error of . As discussed in Section 2.1, we will assume that is relaxed: that is, provides exponential tail bounds on the rank error and on the number of inversions experienced by an element, in terms of the parameter . Intuitively, it may help to think of a queue which returns a uniformly random element of the topk at each step as the “canonical” relaxed . See Figure 1 for an illustration. (As discussed in Section 2.1, real schedulers have slightly different properties, which are captured in our framework.) We state and prove two technical lemmas parameterized by .
First, we characterize the probability that, for some edge in the dependency graph where is a predecessor of , experiences an inversion before is processed. We say that vertex experiences an inversion on or above node at some point during the execution if , but some node with label at least is returned by before is processed during the execution.
Lemma 1.
Proof.
We begin by proving a few immediate claims.
Claim 1.
At any time , the probability of removing the element of top rank from is at least .
Proof.
By the rank bound, we have that It therefore follows that . ∎
Let be the first time when experiences an inversion, and let be its rank at that time. Since an element of rank must be chosen at , we have that, for any ,
In particular, , for any constant . That is, has rank at the time where it experiences its first inversion, w.h.p. in . We now wish to bound the number of removals between the point when experiences its first inversion, and the point when
is removed. Let this random variable be
. By Claim 1, the top element is always removed within trials in expectation, and hence we can show thatfor , by using inequality :
(1) 
We can prove that
by using bounds for the negative binomial distribution with success probability at least
, since success in our case is deleting the top element.We now wish to know the probability that for a fixed one of these steps is an inversion experienced by on or above . Node has lesser priority than , chosen uniformly at random. Let be the position of , noting that , for any integer . Fix a step , and pessimistically assume that is at the top of the queue at this time. Note that if is at the top of the queue, then has rank larger than at time step . From the fairness bound we know that . Conditioned on , we have that rank of is always at least until is removed. This allows us to use the rank bound for upper bounding the probability of experiencing inversion on or above at each step and then we can union bound over steps. Hence, we get that the probability that experiences an inversion on or above is at most . Further, bounding over all choices of , we get the upper bound:
(2) 
Finally, bounding over all possible values of and their probabilities, we get that the probability that experiences an inversion on or above during the execution is at most:
(3) 
∎
Furthermore, the above proof directly implies the following corollary:
Corollary 1.
In Appendix A, we also prove a slightly tighter version of Lemma 1 for the case where the implementation of provides the further guarantee of only returning elements from the topk. Note that such a queue is always rank bounded, but is not necessarily fair.
Our second technical lemma quantifies the expected number of priority inversions incurred by an element, , of once ’s dependencies have been processed—that is, the number of times an element of with lower priority than is returned by before is. If a vertex has no predecessor in at some time , we call a root.
Lemma 2.
Proof.
Follows immediately from the fairness provided by . ∎
We stress that these two lemmas quantify the entire contribution of (the randomness of) the relaxation of to the analysis. The major burden of the analysis, particularly for MIS, is instead to manage the interaction between the randomness of the permutation (which is not inherently related to the relaxation of ) and the structure of . Equipped with these lemmas, we are ready to do just that.
3.1 Algorithm 2: The General Case
The following theorem shows that the relaxed queue in Algorithm 2 has essentially no cost for sparse dependency graphs with and still completes in iterations even for dense dependency graphs when . For example, Theorem 1 demonstrates that taskbased problems which are inherently sparse such as Knuth Shuffle and List Contraction [5] incur only negligible “wasted work” when utilizing a relaxed queue with . Furthermore, graph problems with edge dependencies such as greedy vertex coloring incur a cost proportional to the sparsity of the underlying graph. Although the result is not technically challenging, it is tight up to factors of .
Proof.
We will compute the expected number of failed deletes directly as follows: Whenever a failed delete occurs on a node , charge it to the lexicographically first edge, , for which and are both unprocessed and (i.e., with possibly ). Note that (1) such an edge must exist or else a failed delete could not have occurred, (2) the failed delete must represent a priority inversion on , and (3) must be a root (because is lexicographically first). The first time an edge is charged, we call the active edge until is processed. Since is a root for the duration of ’s status as active edge, by Lemma 2, only experiences priority inversions in expectation while is active, which upper bounds the number of failed deletes charged to .
Let be the event that edge ever becomes active. can only occur if experiences an inversion on or above during the execution, which is bounded by by Lemma 1. Thus, the total expected cost of is at most . There are edges so the total cost is as claimed. ∎
Briefly, to see that Theorem 1 is tight (up to factors of ), consider executing a greedy graph coloring problem on a clique. In this case, at any step, only the highest priority node can ever be processed, and for each such node, , it takes delete attempts before is processed. Thus in total, the algorithm runs for iterations.
3.2 Algorithm 4: Maximal Independent Set
The following theorem bounds the number of iterations of Algorithm 4. By contrast to Algorithm 2, we show that using a relaxed queue for computing Maximal Independent Sets on large graphs has essentially no cost at all, even for dense graphs! In particular, Algorithm 4 incurs a relaxation cost with no dependence on the size or structure of , only on the relaxation factor .
Proof.
Denote the lexicographically first MIS of with respect to as . We first identify the key edges in the execution of Algorithm 4. We will say an edge is a hot edge w.r.t. if is the smallest labeled neighbor of in . Note that if is a hot edge, is not in and has a smaller label than . Let be the event that is a hot edge w.r.t. . Importantly, depends only on the randomness of and not on the randomness of the relaxation of . We make two key observations about hot edges that will allow us to prove the theorem:
Claim 2.
There is exactly one hot edge incident to each vertex , and therefore the total number of hot edges is strictly less than .
This is clear from the condition that is the smallest labeled neighbor of in and the fact that if is not in the , must have at least one neighbor in , or else isn’t maximal.
Claim 3.
A node is only reinserted by Algorithm 4 if there is at least one hot edge with a root and (with possibly ). If is such an edge, we say is active. Furthermore, at least one active hot edge satisfies .
If is reinserted, then must be live and adjacent to some smaller labeled live vertex . Either is a root, in which case is the claimed hot edge, or else must be adjacent to an even smaller labeled live vertex. In the latter case, we can recurse the argument down to and eventually find a hot edge. In either case, both nodes incident to the discovered active hot edge will have a label no greater than ’s.
Proof Outline.
The strategy from here is a follows: whenever a failed delete occurs on a node , we will charge it to an arbitrary hot edge with a root and (of which there must be at least one by Claim 3). Similar to Theorem 1, we will say that is active during the interval between the first time experiences an inversion on or above and the time is processed. We say that the cost, , of an edge, , is the number of failed deletes charged to it (which is notably unless is both hot and, at some point, active). We then separately bound (1) the expected number of active hot edges which ever exist over the execution of Algorithm 4 and (2) the expected number of failed deletes charged to an edge, given that it is an active hot edge. Combining these will give the result.
In order to quantify the distribution of hot edges, we will need one more definition. Fix and let be the subgraph of induced by and let be restricted to . Let be the event that neither nor has a neighbor with . Informally, is the event that both and are still live in after running Algorithm 4 with an exact queue () for iterations but with excluded from . Like , depends only on and not on the randomness of the relaxation of ; furthermore, is independent from and . Using this definition, we can compute:
and
Next, we use the above formulations to bound the probability that a hot edge is ever active. Suppose we are given that is a hot edge and . Then becomes active if and only if suffers an inversion on or above before is processed by the algorithm. Let be the event that becomes active. At this point, we might wish to apply Lemma 1 directly, but unfortunately it is not clear that is independent from , which we will need. However, note that entails but given only that, is otherwise independent from . Thus, if we condition on and , then is fixed and is (conditionally) independent from , and therefore also becomes (conditionally) independent from . Now we can apply Corollary 1, giving
Then:
Observe that for fixed , is decreasing in . In particular, for any permutation in which the event occurs, occurs also, but the reverse is not true. Let . Using Chebyshev’s sum inequality, we obtain:
Finally, since is a root and we only charge for failed deletes on nodes with a larger label than , and therefore a larger label than as well, the number of times we charge is upper bounded by the total number of priority inversions suffered by while a root, which, by Lemma 2, is given by in expectation. Thus .
Combining all the parts, we have a final bound on the total cost:
4 Experimental Results
Synthetic Tests.
To validate our analysis, we implemented the sequential relaxed framework described in Algorithm 2, and used it to solve instances of MIS, matching, Knuth Shuffle, and List Contraction using a relaxed scheduler which uses the MultiQueue algorithm [21], for various relaxation factors. We record the average number of extra relaxations, that is, the number of failed deletes during the entire execution, across five runs. Results are presented in Table 1. We considered graphs of various densities with and vertices. The results appear to confirm our analysis: the number of extra iterations required for MIS is low, and scales only in and not in . There is some variation for fixed and varying , but it is always within a factor of for our trials and does not appear to be obviously correlated with .
4  8  16  32  64  

1000  10000  12.8  56.8  148.8  308.6  583.0 
30000  7.0  40.8  108.6  264.2  478.6  
100000  12.4  40.0  100.6  225.8  427.2  
10000  10000  11.0  43.2  145.4  336.4  738.6 
30000  16.6  71.4  196.0  437.6  890.2  
100000  13.0  56.2  144.4  290.6  529.6 
Concurrent Experiments.
We implemented a simple version of our scheduling framework, using a variant of the MultiQueue [21] relaxed priority queue data structure. We assume a setting where the input, that is, the set of tasks, is loaded initially into the scheduler, and is removed by concurrent threads. We use lockfree lists to maintain the individual priority queues and we hold pointers to the adjacency lists of each node within the queue elements, in order to be able to efficiently check whether a task still has outstanding dependencies.
We compared to the exact scheduling framework using the Waitfree Queue as Fast as FetchandAdd [27]. Since there could still be some reordering of tasks due to concurrency, we elect to use a backoff scheme wherein if an unprocessed predecessor is encountered, we wait for the predecessor to process. In practice this rarely occurs.
Setup.
Our experiments were run on an Intel Xeon Gold 6150 machine with 4 sockets, 18 cores per socket and 2 hyperthreads per core, for a total of 36 threads per socket, and 144 threads total. The machine has 512GB of memory, with 128GB local to each socket. Accesses to local memory are cheaper than accesses to remote memory. We pinned threads to avoid unnecessary context switches and to fill up sockets one at a time. The machine runs Ubuntu 14.04 LTS. We used the GNU C++ compiler (G++) 6.3.0 with optimization level O3.
Experiments are performed on random graphs in three classes; sparse graphs with nodes and edges, small dense graphs with nodes and edges, and large dense graphs with nodes and edges. Our experiments were bottlenecked by graph generation and loading time so we were limited to these graph sizes.
For each data point, we run five trials. In each trial, a graph is generated in parallel by 144 threads, and then we measure the time for threads to compute an MIS. Note that, even when , we generate the graph using 144 threads. To ensure that this does not change memory locality in a way that would invalidate our results, we used the numactl utility to cause memory to be allocated locally on only the sockets where threads will compute the MIS. We verified that this yields the same behavior as experiments where the graph is generated with only threads. (Without using numactl, we observed significant slowdowns in the sequential algorithm.)
The number of queues in the MultiQueue is the number of threads. In our graphs, we plot the average run time on a logarithmic yaxis versus the number of concurrent threads. Error bars show minimum and maximum run times.
Discussion.
Figure 2 shows that our framework using a relaxed scheduler scales with respect to the time to compute MIS over the target graph all the way up to max thread count. The exact framework using the fast waitfree queue also scales, but not as well. In the sparse graphs, the relaxed scheduler is up to faster than optimized sequential code, compared to the exact scheduler, which peaks at faster. In the small dense graphs, where the time spent traversing edges in the MIS algorithm dominates the minor cost of dequeuing nodes, the exact scheduler achieves a peak speedup of over the sequential algorithm, which is approaching the relaxed scheduler’s peak speedup of . However, in the large dense graphs, even though many edges are still traversed by the MIS algorithm, there are sufficiently many nodes to be dequeued that the performance advantage of the relaxed scheduler shows through: it achieves speedup of up to compared to the exact scheduler, which manages only . Note that the single threaded performance of the relaxed scheduler is also quite close to the sequential algorithm. In contrast, the exact scheduler is orders of magnitude slower with a single thread.
5 Future Work
From a theoretical perspective, the natural next step would be to tighten the bound on failed deletes, both for the generic algorithm and for MIS; in fact, we conjecture that the bounds in both Theorems 1 and 2 can be replaced with . However, proving such a bound seems to require a deep understanding of the interplay between the structure of and the effects of the randomness of a relaxed queue, which we had to take care in our analysis to keep separate. Also of interest is to discover more applications, and perhaps more instances like MIS in which the bound in Theorem 1 can be improved on.
One shortcoming of our approach is the fact that our cost measure is the number of vertex accesses in the priority queue. Notice that in theory our bounds may be substantially different when expressed in other metrics, such as the number of edge accesses for the algorithm to terminate, which are closer to standard work bounds. We plan to investigate such cost measures in future work.
From the practical perspective, the immediate step would to improve upon our preliminary results, and implement a highperformance variant of this scheduler, and use this framework in the context of more general graph processing packages.
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Appendix A A Tighter Version of Lemma 1.
We now prove a slightly tighter version for the case where only elements from the topk may be chosen by the scheduler . Note that the condition in the following Lemma (that and are simultaneously in the topk) is a prerequisite for to experience an inversion on or above , and thus the Lemma is slightly stronger than necessary.
Lemma 3.
Proof.
We will write as shorthand for the event that and are simultaneously in the topk of at some time. Note that no matter what dependencies exist in the topk of , the entire topk is flushed after the rank element gets deleted times. The number of iterations it takes to delete the rank element times after enters the topk (thereby flushing with certainty) is a negative binomially distributed random variable with mean and success probability (due to the fairness of ), and similarly for . Since entails that either or , we note that the two cases are symmetric and compute:
(*)  
where (*) uses a standard tail bound on the Negative Binomial Distribution^{2}^{2}2See [16] for a derivation..∎
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