latency-tolerant software distributed shared memory...relational query processing engine. comparing...

This paper is included in the Proceedings of the 2015 USENIX Annual Technical Conference (USENIC ATC ’15).

July 8–10, 2015 • Santa Clara, CA, USA

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Latency-Tolerant Software Distributed Shared Memory

Jacob Nelson, Brandon Holt, Brandon Myers, Preston Brigg, Luis Ceze, Simon Kahan, and Mark Oskin, University of Washington

https://www.usenix.org/conference/atc15/technical-session/presentation/nelson

USENIX Association 2015 USENIX Annual Technical Conference 291

Latency-Tolerant Software Distributed Shared MemoryJacob Nelson, Brandon Holt, Brandon Myers, Preston Briggs, Luis Ceze, Simon Kahan, Mark Oskin

University of WashingtonDepartment of Computer Science and Engineering

AbstractWe present Grappa, a modern take on software distributedshared memory (DSM) for in-memory data-intensive ap-plications. Grappa enables users to program a cluster asif it were a single, large, non-uniform memory access(NUMA) machine. Performance scales up even for ap-plications that have poor locality and input-dependentload distribution. Grappa addresses deficiencies of previ-ous DSM systems by exploiting application parallelism,trading off latency for throughput. We evaluate Grappawith an in-memory MapReduce framework (10⇥ fasterthan Spark [74]); a vertex-centric framework inspired byGraphLab (1.33⇥ faster than native GraphLab [48]); anda relational query execution engine (12.5⇥ faster thanShark [31]). All these frameworks required only 60-690lines of Grappa code.

1 IntroductionData-intensive applications (e.g., ad placement, socialnetwork analysis, PageRank, etc.) make up an importantclass of large-scale computations. Typical hardware com-puting infrastructures for these applications are a collec-tion of multicore nodes connected via a high-bandwidthcommodity network (a.k.a. a cluster). Scaling up per-formance requires careful partitioning of data and com-putation; i.e., programmers have to reason about dataplacement and parallelism explicitly, and for some ap-plications, such as graph analytics, partitioning is diffi-cult. This has led to a diverse ecosystem of frameworks—MapReduce [26], Dryad [43], and Spark [74] for data-parallel applications, GraphLab [48] for certain graph-based applications, Shark [31] for relational queries, etc.They ease development by specializing to algorithmicstructure and dynamic behavior; however, applicationsthat do not fit well into one particular model suffer inperformance.

Software distributed shared memory (DSM) systemsprovide shared memory abstractions for clusters. Histor-ically, these systems [15, 19, 45, 47] performed poorly,largely due to limited inter-node bandwidth, high inter-node latency, and the design decision of piggybacking onthe virtual memory system for seamless global memoryaccesses. Past software DSM systems were largely in-spired by symmetric multiprocessors (SMPs), attemptingto scale that programming mindset to a cluster. How-ever, applications were only suitable for them if theyexhibited significant locality, limited sharing and coarse-grain synchronization—a poor fit for many modern data-

analytics applications. Recently there has been a re-newed interest in DSM research [27, 51], sparked by thewidespread availability of high-bandwidth low-latencynetworks with remote memory access (RDMA) capabil-ity.

In this paper we describe Grappa, a software DSM sys-tem for commodity clusters designed for data-intensiveapplications. Grappa is inspired by the Tera MTA [10,11],a custom hardware-based system. Like the MTA, insteadof relying on locality to reduce the cost of memory ac-cesses, Grappa depends on parallelism to keep processorresources busy and hide the high cost of inter-node com-munication. Grappa also adopts the shared-memory, fine-grained parallel programming mindset from the MTA. Tosupport fine-grained messaging like the MTA, Grappa in-cludes an overlay network that combines small messagestogether into larger physical network packets, therebymaximizing the available bisection bandwidth of com-modity networks. This communication layer is built inuser-space, utilizing modern programming language fea-tures to provide the global address space abstraction. Effi-ciencies come from supporting sharing at a finer granu-larity than a page, avoiding the page-fault trap overhead,and enabling compiler optimizations on global memoryaccesses.

The runtime system is implemented in C++ for a clus-ter of x86 machines with an InfiniBand interconnect, andconsists of three main components: a global address space(§3.1), lightweight user-level tasking (§3.2), and an ag-gregating communication layer (§3.3). We demonstratethe generality and performance of Grappa as a commonruntime by implementing three domain-specific platformson top of it: a simple in-memory MapReduce frame-work; a vertex-centric API (i.e. like GraphLab); and arelational query processing engine. Comparing againstGraphLab itself, we find that a simple, randomly parti-tioned graph representation on Grappa performs 2.5⇥better than GraphLab’s random partitioning and 1.33⇥better than their best partitioning strategy, and scales com-parably out to 128 cluster nodes. The query engine builton Grappa, on the other hand, performs 12.5⇥ faster thanShark on a standard benchmark suite. The flexibility andefficiency of the Grappa shared-memory programmingmodel allows these frameworks to co-exist in the same ap-plication and to exploit application-specific optimizationsthat do not fit within any existing model.

The next section provides an overview of how data-intensive application frameworks can easily and effi-

292 2015 USENIX Annual Technical Conference USENIX Association

ciently map to a shared-memory programming model.§3 describes the Grappa system. §4 presents a quantitiveevaluation of the Grappa runtime. §5 describes relatedwork, and §6 concludes.

2 Data-Intensive Application FrameworksAnalytics frameworks—such as MapReduce, graph pro-cessing and relational query execution—are typically im-plemented for distributed private memory systems (clus-ters) to achieve scale-out performance. While implement-ing these frameworks in a shared-memory system wouldbe straightforward, this has generally been avoided be-cause of scalability concerns. We argue that moderndata-intensive applications have properties that can beexploited to make these frameworks run efficiently andscale well on distributed shared memory systems.

Figure 1 shows a minimal example of implementing a“word count”-like application in actual Grappa DSM code.The input array, chars, and output hash table, cells, aredistributed over multiple nodes. A parallel loop over theinput array runs on all nodes, hashing each key to itscell and incrementing the corresponding count atomically.The syntax and details will be discussed in later sections,but the important thing to note is that it looks similar toplain shared-memory code, yet spans multiple nodes and,as we will demonstrate in later sections, scales efficiently.

Here we describe how three data-intensive computingframeworks map to a DSM, followed by a discussionof the challenges and opportunities they provide for anefficient implementation:

MapReduce. Data parallel operations like map andreduce are simple to think of in terms of shared memory.Map is simply a parallel loop over the input (an array orother distributed data structure). It produces intermediateresults into a hash table similar to that in Figure 1. Reduceis a parallel loop over all the keys in the hash table.

Vertex-centric. GraphLab/PowerGraph is an exam-ple of a vertex-centric execution model, designed for im-plementing machine-learning and graph-based applica-tions [35, 48]. Its three-phase gather-apply-scatter (GAS)API for vertex programs enables several optimizationspertinent to natural graphs. Such graphs are difficult topartition well, so algorithms traversing them exhibit poorlocality. Each phase can be implemented as a parallelloop over vertices, but fetching each vertex’s neighborsresults in many fine-grained data requests.

Relational query execution. Decision support, oftenin the form of relational queries, is an important domainof data-intensive workloads. All data is kept in hashtables stored in a DSM. Communication is a function ofinserting into and looking up in hash tables. One parallelloop builds a hash table, followed by a second parallelloop that filters and probes the hash table, producing

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Figure 2: Grappa design overview

the results. These steps rely heavily on consistent, fine-grained updates to hash tables.The key challenges in implementing these frameworks ona DSM are:

Small messages. Programs written to a shared mem-ory model tend to access small pieces of data, whichwhen executing on a DSM system lead to small inter-nodemessages. What were load or store operations becomecomplex transactions involving small messages over thenetwork. Conversely, programs written using a messagepassing library, such as MPI, expose this complexity toprogrammers, and hence encourage them to optimize it.

Poor locality. As previously mentioned, data-intensiveapplications often exhibit poor locality. For example,how much communication GraphLab’s gather and scatteroperations conduct is a function of the graph partition.Complex graphs frustrate even the most advanced parti-tioning schemes [35]. This leads to poor spatial locality.Moreover, which vertices are accessed varies from itera-tion to iteration. This leads to poor temporal locality.

Need for fine-grain synchronization. Typical data-parallel applications offer coarse-grained concurrencywith infrequent synchronization—e.g., between phasesof processing a large chunk of data. Conversely, graph-parallel applications exhibit fine-grain concurrency withfrequent synchronization—e.g., when done processingwork associated with a single vertex. Therefore, for aDSM solution to be general, it needs to support fine-grainsynchronization efficiently.

Fortunately, data-intensive applications have propertiesthat can be exploited to make DSMs efficient: their abun-dant data parallelism enables high degrees of concurrency;and their performance depends not on the latency of exe-cution of any specific parallel task/thread, as it would infor example a web server, but rather on the aggregate exe-cution time (i.e., throughput) of all tasks/threads. In thenext section we explore how these application propertiescan be exploited to implement an efficient DSM.

3 Grappa DesignFigure 2 shows an overview of Grappa’s DSM system. Be-fore describing the Grappa system in detail, we describeits three main components:

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// distributed input arrayGlobalAddress<char> chars = load_input();

// distributed hash table:using Cell = std::map<char,int>;GlobalAddress<Cell> cells = global_alloc<Cell>(ncells);

forall(chars, nchars, [=](char& c) { // hash the char to determine destination size_t idx = hash(c) % ncells; delegate(&cells[idx], [=](Cell& cell) { // runs atomically if (cell.count(c) == 0) cell[c] = 1; else cell[c] += 1; });});

hash("i")

Figure 1: “Character count” with a simple hash table implemented using Grappa’s distributed shared memory.Distributed shared memory. The DSM system pro-

vides fine-grain access to data anywhere in the system.Every piece of global memory is owned by a particularcore in the system. Access to data on remote nodes isprovided by delegate operations that run on the owningcore. Delegate operations may include normal memoryoperations such as read and write as well as synchronizingoperations such as fetch-and-add [36]. Due to delegation,the memory model offered is similar to what underpinsC/C++ [17, 44], so it is familiar to programmers.

Tasking system. The tasking system supportslightweight multithreading and global distributed work-stealing—tasks can be stolen from any node in the system,which provides automated load balancing. Concurrencyis expressed through cooperatively-scheduled user-levelthreads. Threads that perform long-latency operations(i.e., remote memory access) automatically suspend whilethe operation is executing and wake up when the operationcompletes.

Communication layer. The main goal of our commu-nication layer is to aggregate small messages into largeones. This process is invisible to the application pro-grammer. Its interface is based on active messages [69].Since aggregation and deaggregation of messages needsto be very efficient, we perform the process in parallel andcarefully use lock-free synchronization operations. Forportability, we use MPI [50] as the underlying messaginglibrary as well as for process setup and tear down.

3.1 Distributed Shared MemoryBelow we describe how Grappa implements a sharedglobal address space and the consistency model it offers.

3.1.1 Addressing Modes

Local memory addressing. Applications written forGrappa may address memory in two ways: locally andglobally. Local memory is local to a single core within anode in the system. Accesses occur through conventionalpointers. Applications use local accesses for a numberof things in Grappa: the stack associated with a task, ac-cesses to global memory from the memory’s home core,

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Figure 3: Using global addressing for graph layout.

and accesses to debugging infrastructure local to eachsystem node. Local pointers cannot access memory onother cores, and are valid only on their home core.

Global memory addressing. Grappa allows any localdata on a core’s stacks or heap to be exported to the globaladdress space to be made accessible to other cores acrossthe system. This uses a traditional PGAS (partitionedglobal address space [30]) addressing model, where eachaddress is a tuple of a rank in the job (or global processID) and an address in that process.

Grappa also supports symmetric allocations, which al-locates space for a copy (or proxy) of an object on everycore in the system. The behavior is identical to perform-ing a local allocation on all cores, but the local addressesof all the allocations are guaranteed to be identical. Sym-metric objects are often treated as a proxy to a globalobject, holding local copies of constant data, or allowingoperations to be transparently buffered. A separate pub-lication [41] describes how this was used to implementGrappa’s synchronized global data structures, includingvector and hash map.

Putting it all together. Figure 3 shows an example ofhow global, local and symmetric heaps can all be usedtogether for a simple graph data structure. In this example,vertices are allocated from the global heap, automaticallydistributing them across nodes. Symmetric pointers areused to access local objects which hold information aboutthe graph, such as the base pointer to the vertices, fromany core without communication. Finally, each vertexholds a vector of edges allocated from their core’s localheap, which other cores can access by going through thevertex.

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GlobalAddress<int> A = global_alloc<int>(N);

forall(0, N, [A](int i) { int j = random(i) % N; delegate( A + j, [](int& A_j){ A_j += 1; });});

[](int& A_j){ A_j += 1;}

move execution

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Figure 4: Grappa delegate example.3.1.2 Delegate Operations

Access to Grappa’s distributed shared memory is providedthrough delegate operations, which are short operationsperformed at the memory location’s home node. Whenthe data access pattern has low locality, it is more efficientto modify the data on its home core rather than bringing acopy to the requesting core and returning a modified ver-sion. Delegate operations [49, 53] provide this capability.While delegates can trivially implement read/write oper-ations to global memory, they can also implement morecomplex read-modify-write and synchronization opera-tions (e.g., fetch-and-add, mutex acquire, queue insert).Figure 4 shows an example.

Delegate operations must be expressed explicitly tothe Grappa runtime, a change from the traditional DSMmodel. In practice, even programmers using implicitDSMs had to work to express and exploit locality to ob-tain performance. In other work we have developed acompiler [40] that automatically identifies and extractsproductive delegate operations from ordinary code.

A delegate operation can execute arbitrary code pro-vided it does not lead to a context switch. This guaranteesatomicity for all delegate operations. To avoid contextswitches, a delegate must only touch memory owned bya single core. A delegate is always executed at the homecore of the data addresses it touches. Given these restric-tions, we can ensure that delegate operations for the sameaddress from multiple requesters are always serializedthrough a single core in the system, providing atomic-ity with strong isolation. A side benefit is that atomicoperations on data that are highly contended are faster.When programmers want to operate on data structuresspread across multiple nodes, accesses must be expressedas multiple delegate operations along with appropriatesynchronization operations.

3.1.3 Memory Consistency Model

Accessing global memory though delegate operations al-lows us to provide a familiar memory model. All synchro-nization is done via delegate operations. Since delegateoperations execute on the home core of their operand insome serial order and only touch data owned by that singlecore, they are guaranteed to be globally linearizable [38],with their updates visible to all cores across the system

in the same order. In addition, only one synchronousdelegate will be in flight at a time from a particular task,i.e., synchronization operations from a particular task arenot subject to reordering. Moreover, once one core isable to see an update from a synchronous delegate, allother cores are too. Consequently, all synchronizationoperations execute in program order and are made visi-ble in the same order to all cores in the system. Theseproperties are sufficient to guarantee a memory modelthat offers sequential consistency for data-race-free pro-grams [5], which is what underpins C/C++ [17, 44]. Thesynchronous property of delegates provides a clean modelbut is restrictive: we discuss asynchronous operationswithin the next section.

3.2 Tasking SystemEach hardware core has a single operating system threadpinned to it; all Grappa code runs in these threads. Thebasic unit of execution in Grappa is a task. When a task isready to execute, it is mapped to a user-level worker threadthat is scheduled within an operating system thread; werefer to these as workers to avoid confusion. Schedulingbetween tasks is carried out entirely in user-mode withoutoperating system intervention.

Tasks. Tasks are specified by a closure (also referred toas a “functor” or “function object” in C++) that holds bothcode to execute and initial state. The closure can be speci-fied with a function pointer and explicit arguments, a C++struct that overloads the parentheses operator, or a C++11lambda construct. These objects, typically small (⇠ 32bytes), hold read-only values such as an iteration indexand pointers to common data or synchronization objects.Task closures can be serialized and transported aroundthe system, and are eventually executed by a worker.

Workers. Workers execute application and system(e.g., communication) tasks. A worker is simply a collec-tion of status bits and a stack, allocated at a particular core.When a task is ready to execute it is assigned to a worker,that executes the task closure on its own stack. Once atask is mapped to a worker it stays with that worker untilit finishes.

Scheduling. During execution, a worker yields controlof its core whenever performing a long-latency operation,allowing the processor to remain busy while waiting forthe operation to complete. In addition, a programmercan direct scheduling explicitly. To minimize context-switch overhead, the Grappa scheduler operates entirelyin user-space and does little more than store state of oneworker and load that of another. When a task encountersa long-latency operation, its worker is suspended andsubsequently woken when the operation completes.

Each core in a Grappa system has its own independentscheduler. The scheduler has a collection of active work-ers ready to execute called the ready worker queue. Each

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scheduler also has three queues of tasks waiting to beassigned a worker. The first two run user tasks: a publicqueue of tasks that are not bound to a core yet, and aprivate queue of tasks already bound to the core where thedata they touch is located. The third is a priority queuescheduled according to task-specific deadline constraints;this queue manages high priority system tasks, such asperiodically servicing communication requests.

Context switching. Grappa context switches betweenworkers non-preemptively. As with other cooperativemultithreading systems, we treat context switches as func-tion calls, saving and restoring only the callee-saved stateas specified in the x86-64 ABI [12] rather than the fullregister set required for a preemptive context switch. Thisrequires 62 bytes of storage.

Grappa’s scheduler is designed to support a very largenumber of concurrently-active workers—so large, in fact,that their combined context data will not fit in cache. In or-der to minimize unnecessary cache misses on context data,the scheduler explicitly manages the movement of contextdata into the cache. To accomplish this, we establish apipeline of ready worker references in the scheduler. Thispipeline consists of ready-unscheduled, ready-scheduled,and ready-resident stages. When context prefetching ison, the scheduler is only ever allowed to run workers thatare ready-resident; all other workers are assumed to beout-of-cache. The examined part of the ready queue itselfmust also be in cache. In a FIFO schedule, the head of thequeue will always be in cache due to its spatial locality.Other schedules are possible as long as the amount of datathey need to examine to make a decision is independentof the total number of workers.

When a worker is signaled, its reference is markedready-unscheduled. Every time the scheduler runs, one ofits responsibilities is to pick a ready-unscheduled workerto transition to ready-scheduled: it issues a softwareprefetch to start moving the task toward L1. A workerneeds its metadata (one cache line) and its private workingset. Determining the exact working set might be difficult,but we find that approximating the working set with thetop 2-3 cache lines of the stack is the best naive heuristic.The worker data is ready-resident when it arrives in cache.Since the arrival of a prefetched cache line is generallynot part of the architecture, we must determine the latencyfrom profiling.

At our standard operating point on our cluster (⇡1,000workers), context switch time is on the order of 50 ns. Aswe add workers, the time increases slowly, but levels off:with 500,000 workers context switch time is around 75ns. Without prefetching, context switching is limited bymemory access latency—approximately 120 ns for 1,000workers. Conversely, with prefetching on, context switch-ing rate is limited by memory bandwidth—we determinethis by calculating total data movement based on switch

rate and cache lines per switch in a microbenchmark. As areference point, for the same yield test using kernel-levelPthreads on a single core, the switch time is 450ns for afew threads and 800ns for 1000–32000 threads.

Expressing parallelism. The Grappa API supportsspawning individual tasks, with optional data localityconstraints. These tasks may run as full-fledged workerswith a stack and the ability to block, or they may beasynchronous delegates, which like delegate operationsexecute non-blocking regions of code atomically on asingle core’s memory. Asynchronous delegates are treatedas task spawns in the memory model.

For better programmability, tasks are automaticallygenerated from parallel loop constructs, as in Figure 1.Grappa’s parallel loops spawn tasks using a recursivedecomposition of iterations, similar to Cilk’s cilk for con-struct [16], and TBB’s parallel for [59]. This generatesa logarithmically-deep tree of tasks, stopping to executethe loop body when the number of iterations is below auser-definable threshold.

Grappa loops can iterate over an index space or over aregion of shared memory. In the former case, tasks arespawned with no locality constraints, and may be stolenby any core in the system. In the latter case, tasks arebound to the home core of the piece of memory on whichthey are operating so that the loop body may optimize forthis locality, if available. The local region of memory isstill recursively decomposed so that if a particular loopiteration’s task blocks, other iterations may run concur-rently on the core.

3.3 Communication SupportGrappa’s communication layer has two components: auser-level messaging interface based on active messages,and a network-level transport layer that supports requestaggregation for better communication bandwidth.

Active message interface. At the upper (user-level)layer, Grappa implements asynchronous active mes-sages [69]. Our active messages are simply a C++11lambda or other closure. We take advantage of the factthat our homogeneous cluster hardware runs the samebinary in every process: each message consists of atemplate-generated deserializer pointer, a byte-for-bytecopy of the closure, and an optional data payload.

Message aggregation. Since communication is veryfrequent in Grappa, aggregating and sending messagesefficiently is very important. To achieve that, Grappamakes careful use of caches, prefetching, and lock-freesynchronization operations.

Figure 5 shows the aggregation process. Cores keeptheir own outgoing message lists, with as many entries asthe number of system cores in a Grappa system. Theselists are accessible to all cores in a Grappa node to allowcores to peek at each other’s message lists. When a task

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Figure 5: Message aggregation process.sends a message, it allocates a buffer from a pool, deter-mines the destination system node, writes the messagecontents into the buffer, and links the buffer into the corre-sponding outgoing list. These buffers are referenced onlytwice for each message sent: once when the message iscreated, and (much later) when the message is serializedfor transmission. The pool allocator prefetches the bufferswith the non-temporal flag to minimize cache pollution.

Each processing core in a given system node is respon-sible for aggregating and sending the resulting messagesfrom all cores on that node to a set of destination nodes.Cores periodically execute a system task that examinesthe outgoing message lists for each destination node forwhich the core is responsible; if the list is long enoughor a message has waited past a time-out period, all mes-sages to a given destination system node from that sourcesystem node are sent by copying them to a buffer visibleto the network card. Actual message transmission can bedone purely in user-mode using MPI, which in turn usesRDMA.

The final message assembly process involves manipu-lating several shared data-structures (the message lists),so it uses CAS (compare-and-swap) operations to avoidhigh synchronization costs. This traversal requires carefulprefetching because most of the outbound messages arenot in the processor cache at this time (recall that a corecan be aggregating messages originating from other coresin the same node). Note that we use a per-core array ofmessage lists that is only periodically modified acrossprocessor cores, having experimentally determined thatthis approach is faster (sometimes significantly) than aglobal per-system node array of message lists.

Once the remote system node has received the messagebuffer, a management task is spawned to manage theunpacking process. The management task spawns a taskon each core at the receiving system to simultaneouslyunpack messages destined for that core. Upon completion,these unpacking tasks synchronize with the managementtask. Once all cores have processed the message buffer,the management task sends a reply to the sending systemnode indicating the successful delivery of the messages.

3.3.1 Why not just use native RDMA support?

Given the increasing availability and decreasing cost ofRDMA-enabled network hardware, it would seem log-

ical to use this hardware to implement Grappa’s DSM.Figure 6 shows the performance difference between na-tive RDMA atomic increments and Grappa atomic incre-ments using the GUPS cluster-wide random access bench-mark using the cluster described in §4. The cluster hasMellanox ConnectX-2 40Gb InfiniBand cards connectedthrough a QLogic switch with no oversubscription. TheRDMA setting of the experiment used the network card’snative atomic fetch-and-increment operation, and issuedincrements to the card in batches of 512. The Grappasetting issued delegate increments in a parallel for loop.Both settings perform increments to random locations in a32 GB array of 64-bit integers distributed across the clus-ter. Figure 6(left) shows how aggregation allows Grappato exceed the performance of the card by 25⇥ at 128nodes. We measured the effective bisection bandwidth ofthe cluster as described in [39]: for GUPS, performanceis limited by memory bandwidth during aggregation, anduses ⇠ 40% of available bisection bandwidth.

Figure 6(right) illustrates why using RDMA directlyis not sufficient. The data also shows that MPI over In-finiBand has negligible overhead. Our cluster’s cards areunable to push small messages at line rate into the net-work: we measured the peak RDMA performance of ourcluster’s cards to be 3.2 million 8-byte writes per second,when the wire-rate limit is over 76 million [42]. We be-lieve this limitation is primarily due to the latency of themultiple PCI Express round trips necessary to issue oneoperation; a similar problem was studied in [34]. Fur-thermore, RDMA network cards have severely limitedsupport for synchronization with the CPU [27, 51]. Fi-nally, framing overheads can be large: InfiniBand 8-byteRDMA writes moves 50 bytes on the wire; Ethernet-based RDMA using RoCE moves 98 bytes. Work isongoing to improve network card small message perfor-mance [1, 4, 28, 34, 55, 57, 61, 68]: even if native smallmessage performance improves in future hardware, ouraggregation support will still be useful to minimize cacheline movement, PCI Express round trips, and other mem-ory hierarchy limitations.

3.4 Fault tolerance discussionA number of recent “big data” workload studies [22, 60,62] suggest that over 90 percent of current analytics jobsrequire less than one terabyte of input data and run forless than one hour. We designed Grappa to support thissize of workload on medium-scale clusters, with tens tohundreds of nodes and a few terabytes of main memory.At this scale, the extreme fault tolerance found in systemslike Hadoop is largely wasted — e.g., assuming a per-machine MTBF of 1 year, we would estimate the MTBFof our 128-node cluster to be 2.85 days.

We could add checkpoint/restart functionality toGrappa, either natively or using a standard HPC library

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Figure 6: On the left, random updates to a billion-integerdistributed array with the GUPS benchmark. On the right,ping-pong bandwidth measured between two nodes.

[29]. Writing a checkpoint would take on the order ofminutes; for instance, our cluster can write all 8 TB ofmain memory to its parallel filesystem in approximately10 minutes. It is important to balance the cost of takinga checkpoint with the work lost since the last checkpointin the event of a failure. We can approximate the opti-mum checkpoint interval using [73]; assuming a check-point time of 10 minutes and a per-machine MTBF of 1year, we should take checkpoints every 4.7 hours. Theseestimates are similar to what Low et al. measured forGraphlab’s checkpoint mechanism in [48]. In this regime,it is likely cheaper to restart a failed job than it is to paythe overhead of taking checkpoints and recovering from afailure.

Given these estimates, we chose not to implement faulttolerance in this work. Adding more sophisticated faulttolerance to Grappa for clusters with thousands of nodesis an interesting area of future work.

4 EvaluationWe implemented Grappa in C++ for the Linux operatingsystem. The core runtime system system is 17K linesof code. We ran experiments on a cluster of AMD Inter-lagos processors with 128 nodes. Nodes have 32 coresoperating at 2.1GHz, spread across two sockets, 64GBof memory, and 40Gb Mellanox ConnectX-2 InfiniBandnetwork cards. Nodes are connected via a QLogic In-finiBand switch with no oversubscription. We used astock OS kernel and device drivers. The experimentswere run in a machine without administrator access orspecial privileges. GraphLab and Spark communicatedusing IP-over-InfiniBand in Connected mode.

4.1 Vertex-centric Programs on GrappaWe implemented a vertex-centric programming frame-work in Grappa with most of the same core functionalityas GraphLab [35, 48] using the graph data structure pro-vided by the Grappa library (Figure 3). Unlike GraphLabwe do not focus on intelligent partitioning, instead choos-

ing a simple random placement of vertices to cores. Edgesare stored co-located on the same core with vertex data.Using this graph representation, we implement a subsetof GraphLab’s synchronous engine, including the deltacaching optimization, in ⇠60 lines of Grappa code. Par-allel iterators are defined over the vertex array and overeach vertex’s outgoing edge list. Given our graph struc-ture, we can efficiently support gather on incoming edgesand scatter on outgoing edges. Users of our Vertex-centricGrappa framework specify the gather, apply, and scatteroperations in a “vertex program” structure. Vertex pro-gram state is represented as additional data attached toeach vertex. The synchronous engine consists of severalparallel forall loops executing the gather, apply, andscatter phases within an outer “superstep” loop until allvertices are inactive.

We implemented three graph analytics applicationsfrom GraphBench [3] using vertex program definitionsequivalent to GraphLab’s: PageRank, Single SourceShortest Path (SSSP), and Connected Components (CC).In addition, we implemented a simple Breadth-first search(BFS) application in the spirit of the Graph500 bench-mark [37], which finds a “parent” for each vertex with agiven source. The implementation in the GraphLab APIis similar to the SSSP vertex program.

4.1.1 Performance

To evaluate Grappa’s Vertex-centric framework imple-mentation, we ran each application on the Twitter fol-lower graph [46] (41 M vertices, 1 B directed edges) andthe Friendster social network [72] (65 M vertices, 1.8 Bundirected edges). For each we run to convergence—forPageRank we use GraphLab’s default threshold criteria—resulting in the same number of iterations for each. Addi-tionally, for PageRank we ran with delta caching enabled,as it proved to perform better. For Grappa we use the no-replication graph structure with random vertex placement;for GraphLab, we show results for random partitioningand the current best partitioning strategy: “PDS” whichcomputes the “perfect difference set”, but can only be runwith p2 + p + 1 (where p is prime) nodes. Most of thecomparisons are done at 31 nodes for this reason.

Figure 7a depicts performance results at 31 nodes, nor-malized to Grappa’s execution time. We can see thatGrappa is faster than random partitioning on all the bench-marks (on average 2.57⇥), and 1.33⇥ faster than the bestpartitioning, despite not replicating the graph at all. Bothimplementations of PageRank issue application-level re-quests on the order of 32 bytes (mostly communicatingupdated rank values). However, since these would per-form terribly on the network, both systems aggregateupdates into larger wire-level messages. Grappa’s per-formance exceeds that of GraphLab primarily because itdoes this faster.

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Figure 7c(bottom) explores this difference using theGUPS benchmark from §3.3.1. All systems send 32-byte updates to random nodes which then update a 64-bit word in memory: this experiment models only thecommunication of PageRank and not the activation ofvertices, etc. For GraphLab and Spark, the messaginguses TCP-over-IPoIB and the aggregators make 64KBbatches (GraphLab also uses MPI, but for job startuponly). At 31 nodes, GraphLab’s aggregator achieves 0.14GUPS, while Grappa achieves 0.82 GUPS. Grappa’s useof RDMA accounts for about half of that difference; whenGrappa uses MPI-over-TCP-over-IPoIB it achieves 0.30GUPS. The other half comes from Grappa’s prefetching,more efficient serialization, and other messaging designdecisions. The Spark result is an upper bound obtainedby writing directly to Spark’s java.nio-based messagingAPI rather than Spark’s user-level API.

During the PageRank computation, Grappa’s unsophis-ticated graph representation sends 2⇥ as many messagesas GraphLab’s replicated representation. However, as canbe seen in Figure 7c(top), Grappa sends these messagesat up to 4⇥ the rate of GraphLab over the bulk of its exe-cution. At the end of the execution when the number ofactive vertices is low, both systems’ message rates drop,but Grappa’s simpler graph representation allows it toexecute these iterations faster as well. Overall, this leadsto a 2⇥ speedup.

Figure 8 demonstrates the connection between concur-rency and aggregation over time while executing PageR-ank. We see that at each iteration, the number of concur-rent tasks spikes as scatter delegates are performed on

outgoing edges, which leads to a corresponding spike inbandwidth due to aggregating the many concurrent mes-sages. At these points, Grappa achieves roughly 1.1 GB/sper node, which is 47% of peak bisection bandwidth forlarge packets discussed in §3.3.1, or 61% of the band-width for 80 kB messages, the average aggregated size.This discrepancy is due to not being able to aggregatepackets as fast as the network can send them, but is stillsignificantly better than unaggregated bandwidth.

Figure 7b(left) shows strong scaling results on bothdatasets. As we can see, scaling is poor beyond 32 nodesfor both platforms, due to the relatively small size of thegraphs—there is not enough parallelism for either systemto scale on this hardware. To explore how Grappa fareson larger graphs, we show results of a weak scaling exper-iment in Figure 7b(right). This experiment runs PageR-ank on synthetic graphs generated using Graph500’s Kro-necker generator, scaling the graph size with the numberof nodes, from 200M vertices, 4B edges, up to 2.1B ver-tices, 34B edges. Runtime is normalized to show distancefrom ideal scaling (horizontal line), showing that scalingdeteriorates less than 30% at 128 nodes.

4.2 Relational queries on Grappa

We used Grappa to build a distributed backend to Raco, arelational algebra compiler and optimization framework[58]. Raco supports a variety of relational query languagefrontends, including SQL, Datalog, and an imperativelanguage, MyriaL. It includes an extensible relationalalgebra optimizer and various intermediate query planrepresentations.

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We compare performance of our system to that ofShark, a fast implementation of Hive (SQL-like), builtupon Spark. We chose this comparison point becauseShark is optimized for in-memory execution and performscompetitively with parallel databases [71].

Our particular approach for the Grappa backend toRaco is source-to-source translation. We generateforalls for each pipeline in the physical query plan.We extend the code generation approach for serial codein [54] to generating parallel shared memory code. Thegenerated code is sent through a normal C++11 compiler.

All data structures used in query execution (e.g. hashtables for joins) are globally distributed and shared. Whilethis a departure from the shared-nothing architecture ofnearly all parallel databases, the locality-oriented execu-tion model of Grappa makes the execution of the queryvirtually identical to that of traditional designs. We ex-pect (and later demonstrate) that Grappa will excel at hashjoins, given that it achieves high throughput on randomaccess.

Implementing the parallel Grappa code generation wasa relatively simple extension of the generator for serialC++ code that we use for testing Raco. It required lessthan 90 lines of template C++/Grappa code and 600 linesof support and data structure C++/Grappa code to imple-ment conjunctive queries, including two join implementa-tions.

4.2.1 Performance

We focus on workloads that can be processed in memory,since storage is out of scope for this work. For Grappa, wescan all tables into distributed arrays of rows in memory,then time the query processing. To ensure all timed pro-cessing in Shark is done in memory, we use the methodol-ogy that Shark’s developers use for benchmarking [2]. In

particular, all input tables are cached in memory and theoutput is materialized to an in-memory table. The numberof reducer tasks for shuffles was set to 3 per Spark worker,which balances overhead and load balance. Each workerJVM was assigned 52GB of memory.

We ran conjunctive queries from SP2Bench [63]. Thequeries in this benchmark involve several joins, whichmakes it interesting for evaluating parallel in-memory sys-tems. We show results on 16 nodes (we found Shark failedto scale beyond 16 nodes on this data set) in Figure 9a.Grappa has a geometric mean speedup of 12.5⇥ overShark. The benchmarks vary in performance due to dif-ferences in magnitude of communication and output size.

There are many differences between the two runtimesystems (e.g. messaging layers, JVM and native) and thequery processing approach (e.g. iterators vs compiledcode), making it challenging to clearly understand thesource of the performance difference between the twosystems. To do so, we computed a detailed breakdown(Figure 9b) of the execution of Q2. We took sample-basedprofiles of both systems and categorized CPU time intofive components: network (low-level networking over-heads, such as MPI and TCP/IP messaging), serialization(aggregation in Grappa, Java object serialization in Shark),iteration (loop decomposition and iterator overheads), ap-plication (actual user-level query directives), and other(remaining runtime overheads for each system).

Overall, we find that the systems spend nearly the sameamount of CPU time in application computation, andthat more than half of Grappa’s performance advantagecomes from efficient message aggregation and a moreefficient network stack. An additional benefit comes fromiterating via Grappa’s compiled parallel for-loops com-pared to Shark’s dynamic iterators. Finally, both systemshave other, unique overheads: Grappa’s scheduling timeis higher than Shark due to frequent context switches,whereas Shark spends time dynamically checking thetypes of data values.

Shark’s execution of these queries appears to placebursty demands on the network, and is sensitive to net-work bandwidth. On query Q2, Shark achieves thesame peak bandwidth as GUPS (Figure 7c) sustains(200MB/s/node), but its sustained bandwidth is just overhalf this amount (116 MB/s/node).

4.3 Iterative MapReduce on GrappaWe experiment with data parallel workloads by imple-menting an in-memory MapReduce API in 152 lines ofGrappa code. The implementation involves a forall overinputs followed by a forall over key groups. In the all-to-all communication, mappers push to reducers. As withother MapReduce implementations, a combiner functioncan be specified to reduce communication. In this case,the mappers materialize results into a local hash table, us-

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ing Grappa’s partition-awareness. The global-view modelof Grappa allows iterations to be implemented by theapplication programmer with a while loop.

4.3.1 Performance

We pick k-means clustering as a test workload; it exer-cises all-to-all communication and iteration. To providea reference point, we compare the performance to theSparkKMeans implementation for Spark. Both versionsuse the same algorithm: map the points, reduce the clus-ter means, and broadcast local means. The Spark codecaches the input points in memory and does not persist par-titions. Currently, our implementation of MapReduce isnot fault-tolerant. To ensure the comparison is fair, wemade sure Spark did not use fault-tolerance features: weused MEMORY ONLY storage level for RDDs, which doesnot replicate an RDD or persist it to disk and verifiedduring the runs that no partitions were recomputed due tofailures. We run k-means on a dataset from Seaflow [66],where each instance is a flow cytometry sample of sea-water containing characteristics of phytoplankton cells.The dataset is 8.9GB and contains 123M instances. Theclustering task is to identify species of phytoplankton sothe populations may be counted.

The results are shown in Figure 10 for K = 10and K = 10000. We find Grappa-MapReduce to benearly an order of magnitude faster than the comparableSpark implementation. Absolute runtime for Grappa-MapReduce is 0.13s per iteration for K = 10 and 17.3sper iteration for K = 10000, compared to 1s and 170srespectively for Spark.

We examined profiles to understand this difference.We see similar results as with Shark: the bulk of thedifference comes from the networking layer and fromdata serialization. As K grows, this problem should becompute-bound: most execution time is spent assigningpoints to clusters in the map step. At large K, Grappa-MapReduce is clearly compute-bound, but Spark spendsonly 50% of its time on compute; the rest is in network

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code in the reduce step. Grappa’s efficient small messagesupport and support for overlapping communication andcomputation help it perform well here.

4.4 Writing directly to GrappaNot all problems fit perfectly into current restricted pro-gramming models—for many, a better solution can befound by breaking these restrictions. An advantage ofbuilding specialized systems on top of a flexible, high-performance platform is that it makes it easier to imple-ment new optimizations into domain-specific models, orimplement a new algorithm from scratch natively. For ex-ample, for BFS, Beamer’s direction-optimizing algorithmhas been shown to greatly improve performance on theGraph500 benchmark by traversing the graph “bottom-up”in order to visit a subset of the edges [13]. This cannot bewritten in a pure Vertex-centric framework like GraphLab.We implemented the Beamer’s BFS algorithm directly onthe existing graph data structure in 70 lines of code. Per-formance results in Figure 11 show that this algorithm’s

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performance is nearly a factor of 2 better than the pureVertex-centric abstraction can achieve.

5 Related WorkMultithreading. Hardware-based implementations ofmultithreading to tolerate latency include the DenelcorHEP [65], Tera MTA [11], Cray XMT [33], Simultaneousmultithreading [67], MIT Alewife [6], Cyclops [9], andGPUs [32]. Hardware multithreading often pays withlower single-threaded performance that may limit appealin the mainstream market. As a software implementa-tion of multithreading for mainstream general-purposeprocessors, Grappa provides the benefits of latency toler-ance only when warranted, leaving single-threaded per-formance intact.

Grappa’s closest software-based multithreading ances-tor is the Threaded Abstract Machine (TAM) [24]. TAMis a software runtime system designed for prototypingdataflow execution models on distributed memory super-computers. Like Grappa, TAM supports inter-node com-munication, management of the memory hierarchy, andlightweight asynchronous scheduling of tasks to proces-sors, all in support of computational throughput despitethe high latency of communications. A notable conclu-sion [25] was that threading for latency tolerance wasfundamentally limited because the latency of the top-levelstore (e.g. L1 cache) is in direct competition with thenumber of contexts that can fit in it. However, we findprefetching is effective at hiding DRAM latency in con-text switching. Indeed, a key difference between Grappa’ssupport for lightweight threads and that of other user levelthreading packages, such as QThreads [70], TBB [59],Cilk [16] and Capriccio [14] is Grappa’s context prefetch-ing. Grappa’s prefetching could likely improve from com-piler analyses inspired by those of Capriccio for reducingmemory usage.

Software distributed shared memory. Much of theinnovation in DSM over the past 30 years has focusedon reducing the synchronization costs of updates. Thefirst DSM systems, including IVY [47], used frequentinvalidations to provide sequential consistency, induc-ing high communication costs for write-heavy workloads.Later systems relaxed the consistency model to reducecommunication demands; some systems further mitigatedperformance degradation due to false sharing by adoptingmultiple writer protocols that delay integration of concur-rent writes made to the same page. The Munin [15, 19]and TreadMarks [45] systems exploited both of theseideas, but still incurred some coherence overhead. Muninand Blizzard [64] allowed the tracking of ownership withvariable granularity to reduce the cost due to false shar-ing. Grappa follows the lead of TreadMarks and providesDSM entirely at user-level through a library and runtime.FaRM [27] offers lower latency and higher throughput up-

dates to DSM than TCP/IP via lock free and transactionalaccess protocols exploiting RDMA, but remote accessthroughput is still limited to the RDMA operation ratewhich is typically an order of magnitude less than the pernode network bandwidth.

Partitioned Global Address Space languages. Thehigh-performance computing community has largely dis-carded the coherent distributed shared memory approachin favor of the Partitioned Global Address Space (PGAS)model. Examples include Split-C [23], Chapel [20],X10 [21], Co-array Fortran [56] and UPC [30]. What ismost different between general DSM systems and PGASones is that remote data accesses are explicit, therebyencouraging developers to use them judiciously. Grappafollows this approach, implementing a PGAS system atthe language level, thereby facilitating compiler and pro-grammer optimizations.

Distributed data-intensive processing frameworks.There are many other data-parallel frameworks likeHadoop, Haloop [18], and Dryad [43]. These are de-signed to make parallel programming on distributed sys-tems easier; they meet this goal by targeting data-parallelprograms. There have also been recent efforts to buildparameter servers for distributed machine learning al-gorithms using asynchronous communication and dis-tributed key-value storage built from RPCs [7, 8]. Theincremental data-parallel system Naiad [52] achieves bothhigh-throughput for batch workloads and low-latency forincremental updates. Most of these designs eschew DSMas an application programming model for performancereasons.

6 Conclusions

Our work builds on the premise that writing data-intensiveapplications and frameworks in a shared memory envi-ronment is simpler than developing custom infrastructurefrom scratch. To that end, Grappa is inspired not by SMPsystems, but by novel supercomputer hardware – the CrayMTA and XMT line of machines. This work borrowsthe core insight of those hardware systems and builds itinto a software runtime tuned to extract performance fromcommodity processors, memory systems and networks.Based on this premise, we show that a DSM system canbe efficient for this application space by judiciously ex-ploiting the key application characteristics of concurrencyand latency tolerance. Our data demonstrates that frame-works such as MapReduce, vertex-centric computation,and query execution are easy to build and efficient. OurMapReduce and query execution implementations are anorder of magnitude faster than the custom frameworksfor each. Our vertex-centric GraphLab-inspired API is1.33⇥ faster than GraphLab itself, without the need forcomplex graph partitioning schemes.

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AcknowledgementsThis work was supported by NSF Grant CCF-1335466,Pacific Northwest National Laboratory, and gifts fromNetApp and Oracle.

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