Correct and Efficient Work-Stealing for Weak Memory Models

13 Pages·2017·1.16 MB·English

Correct and Efficient Work-Stealing for Weak Memory
Models
Nhat Minh Lê, Antoniu Pop, Albert Cohen, Francesco Zappa Nardelli
To cite this version:
Nhat Minh Lê, Antoniu Pop, Albert Cohen, Francesco Zappa Nardelli. Correct and Efficient Work-
Stealing for Weak Memory Models. PPoPP ’13 - Proceedings of the 18th ACM SIGPLAN sym-
posium on Principles and practice of parallel programming, Feb 2013, Shenzhen, China. pp.69-80,
￿10.1145/2442516.2442524￿. ￿hal-00802885￿
HAL Id: hal-00802885
https://hal.inria.fr/hal-00802885
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Correct and Efficient Work-Stealing for Weak Memory Models
NhatMinhLê AntoniuPop AlbertCohen FrancescoZappaNardelli
INRIAandENSParis
Abstract Lev’sdequefortheARMarchitectures,andproveitscorrectness
againstthememorysemanticsdefinedin[12]and[7].Oursecond
ChaseandLev’sconcurrentdequeisakeydatastructureinshared-
contribution is a systematic study of the performance of several
memoryparallelprogrammingandplaysanessentialroleinwork-
implementationsofChase–Levonrelaxedhardware.Indetail,we
stealing schedulers. We provide the first correctness proof of an
compare our optimized ARM implementation against a standard
optimizedimplementationofChaseandLev’sdequeontopofthe
implementationforthex86architectureandtwoportablevariants
POWERandARMarchitectures:theseprovideveryrelaxedmem-
expressed in C11: a reference sequentially consistent translation
orymodels,whichweexploittoimproveperformancebutconsider-
of the algorithm, and an aggressively optimized version making
ablycomplicatethereasoning.Wealsostudyanoptimizedx86and
full use of the release–acquire and relaxed semantics offered by
aportableC11implementation,conductingsystematicexperiments
C11 low-levelatomics. These implementationsof the Chase–Lev
toevaluatetheimpactofmemorybarrieroptimizations.Ourresults
dequeareevaluatedinthecontextofawork-stealingscheduler.We
demonstratethebenefitsofhandtuningthedequecodewhenrun-
considerdiverseworker/thiefconfigurations,includingasynthetic
ningontopofrelaxedmemorymodels.
benchmarkwithtwodifferentworkloadsandstandardtask-parallel
Categories and Subject Descriptors D.1.3 [Programming Tech- kernels. Our experiments demonstrate the impact of the memory
niques]: Concurrent Programming; E.1 [Data Structures]: Lists, barrier optimization on the throughput of our work-stealing run-
stacks,andqueues time. We also comment on how the ARM correctness proof can
be tailored to these alternative implementations. As a side effect,
Keywords lock-free algorithm, work-stealing, relaxed memory
we highlight that our optimized ARM implementation cannot be
model,proof
expressed using C11 low-level atomics, which invariably end up
insertingoneredundantsynchronizationinstruction.
1. Introduction
MulticorePOWERandARMarchitecturesarestandardtargetsfor 2. Chase–Levdeque
server, consumer electronics, and embedded control applications.
User-space runtime schedulers offer an excellent playground for
Thedifficultiesofparallelprogrammingareexacerbatedbythere-
studyinglow-levelhigh-performancecode.Wefocusonrandom-
laxed memory model implemented by these architectures, which
ized work-stealing: it was originally designed as the scheduler
allowtheprocessorstoperformawiderangeofoptimizations,in-
of the Cilk language for shared-memory multiprocessors [4], but
cludingthread-localreorderingandnon-atomicstorepropagation.
thanks to its merits [2] it has been adopted in a number of par-
The safety-critical nature of many embedded applications call
allel libraries and parallel programming environments, including
forsolidfoundationsforparallelprogramming.Thispapershows
theIntelTBBandcompilersuite.Work-stealingvariantshavealso
thatahighdegreeofconfidencecanbeachievedforhighlyopti-
beenproposedfordistributedclusters[5]andheterogeneousplat-
mized,real-world,concurrentalgorithms,runningontopofweak
forms[1].Theschedulingstrategyisintuitive:
memory models. A good test-case is provided by the runtime
• Eachprocessorusesadynamicarrayasadequeholdingtasks
schedulerofatasklibrary.WethusfocusontheChaseandLev’s
readytobescheduled.
concurrent doubly-ended queue (deque) [3], the cornerstone of
• Eachprocessormanagesitsowndequeasastack.Itmayonly
mostwork-stealingschedulers.Untilnow,norigorouscorrectness
pushandpoptasksfromthebottomofitsowndeque.
proofhasbeenbeenprovidedforimplementationsofthisalgorithm
• Otherprocessorsmaynotpushorpopfromthatdeque;instead,
running on top of a relaxed memory model. Furthermore, while
theystealtasksfromthetopwhentheirowndequeisempty.In
work-stealingiswidelyusedonthex86architecture(anevaluation
mostimplementations,thestolendequeisselectedatrandom.
underarestrictivehypothesisofidempotenceoftheworkloadcan
• Initially,oneprocessorstartswiththe“root”taskoftheparallel
befoundin[10]),fewexperimentstargetweakermemorymodels.
programinitsdeque,andallotherdequesareempty.
Ourfirstcontributionisacorrectnessproofofthisfundamen-
Thestate-of-the-artalgorithmforthework-stealingdequeisChase
talconcurrentdatastructurerunningontopofarelaxedmemory
and Lev’s lock-free deque [3]. It uses an array with automatic,
model.Weprovideahand-tunedimplementationoftheChaseand
asynchronous growth. Assuming sequentially consistent memory,
it involves only one atomic compare-and-swap (CAS) per steal,
noCASonpush,andnoCASontakeexceptwhenthedequehas
exactlyonlyoneelementleft.
Permissiontomakedigitalorhardcopiesofallorpartofthisworkforpersonalor We implemented and tested four versions of the concurrent
classroomuseisgrantedwithoutfeeprovidedthatcopiesarenotmadeordistributed dequealgorithm,withdifferentbarrierconfigurations:(1)asequen-
forprofitorcommercialadvantageandthatcopiesbearthisnoticeandthefullcitation
onthefirstpage.Tocopyotherwise,torepublish,topostonserversortoredistribute tiallyconsistentversion,writtenwithC11seq_cstatomics,follow-
tolists,requirespriorspecificpermissionand/orafee. ingtheoriginaldescriptionin[3];(2)anoptimizedversion,which
PPoPP’13, February23–27,2013,Shenzhen,China. takes full advantage of the C11 relaxed memory model, reported
Copyright(cid:13)c 2013ACM978-1-4503-1922/13/02...$10.00 inFigure1;(3)anativeversionforARMv7,reportedinFigure2,
and(4)anativeversionforx86.Thesenativeversionsrelyoncom- rithm.However,iflow-levelatomicsarecompiledusingthemap-
pilerintrinsicsandinlineassemblytoleveragearchitecture-specific pingofMcKenneyandSilvera[9]onARMv7/POWERorthemap-
assumptionsandthusreducethenumberofbarriersrequired. ping of Tehrekov [14] on x86, the generated code contains more
In our implementations of Figure 1 and Figure 2, we assume barriersthanthehand-optimizednativeversionsonbothx86and
thattheDequetypeisdeclaredas: ARMv7. We show in Section 5 that this happens because of the
typedef struct { typedef struct { need for seq_cst atomics to simulate ARMv7/POWER cumula-
aattoommiicc__siinzte_btufsfiezre[;]; aAttoommiicc_(sAirzrea_yt*)toapr,rabyo;ttom; tive semantics. Concretely, on ARMv7, an extra dmb instruction
} Array; } Deque; isinsertedbeforeeachCASoperation[11],comparedtothenative
InthecodeofFigure1theatomic_andmemory_order_prefixes versionwherearelaxedCAS—coherentandatomiconly—issuf-
havebeenelidedforclarity.TheARMv7pseudo-codeofFigure2 ficient. On x86, an mfence instruction is added between the two
usesthekeywordsRandWtodenotereadsandwritestosharedvari- readsinsteal.ThefullysequentiallyconsistentC11implementa-
ables,andatomicindicatesablockthatwillbeexecutedatomically, tioninsertsmanymoreredundantbarriers[11].
implementedviaLL/SCinstructions.Thex86versionisbasedon
priorwork[10]andonlyrequiresasinglemfencememorybarrier 3. ThememorymodelofARMv7
intake,inplaceofthecalltothread_fenceintheC11code.
The memory model of the ARMv7 architecture follows closely
thatofthePOWERarchitecture,allowingawiderangeofrelaxed
2.1 Notionsofcorrectness
behaviorstobeobservabletotheprogrammer:
Theexpectedbehaviorofthework-stealingdequeisintuitive:tasks
pushedintothedequearetheneithertakeninreverseorderbythe 1. The hardware threads can each perform reads and writes out-
samethread,orstolenbyanotherthread.Wesaythatanimplemen- of-order, or even speculatively. Basically any local reordering
tationiscorrectifitsatisfiesfourcriteria, formalized andproven isallowedunlessthereisadata/controldependenceorsynchro-
correctforourARMv7optimizedcodeinSection4: nizationinstructionpreventingit.
1. tasksaretakeninreverseorder; 2. The memory system does not guarantee that a write becomes
2. onlytaskspushedaretakenorstolen(well-definedreads); visible to all other hardware threads at the same time point.
3. ataskpushedintoadequecannotbetakenorstolenmorethan Writesperformedbyonethreadarepropagatedto(andbecome
once(uniqueness); visible from) any other thread in an arbitrary order, unless
4. givenafinitenumberofpushoperations,allpushedvalueswill synchronizationinstructionsareused.
eventuallybeeithertakenorstolenexactlyonce,ifenoughtake 3. A dmb barrier instruction guarantees that all the writes which
andstealoperationsareattempted(existence). havebeenobservedbythethreadissuingthebarrierinstruction
Thesecriteriaholdbecauseofthefollowingassumptionsandprop- are propagated to all the other threads before the thread can
ertiesoftheChase–Levalgorithm: continue.Observedwritesincludeallwritespreviouslyissued
• Foranygivendeque,pushandpopoperationsexecuteonasin- bythethreaditself,aswellasanywritepropagatedtoitfrom
glethread.Concurrencycanonlyoccurbetweenoneexecution another thread prior to the barrier. This semantics of barrier
ofpushortakeintheownerthread,andoneormoreexecutions instructionsisreferredtoascumulative.
ofstealindifferentthreads. WebuildontheaxiomaticformalizationofPOWERandARMv7
• Newly pushed tasks are made visible to take and steal by the memorymodelbyMador-Haimetal.[7],whichhasbeenproved
incrementtobottominpush.AsweshallseeinSection4,our equivalent to the operational semantics of Sarkar et al. [12]. A
ARMv7 implementation enforces this by placing a sync bar- gentleintroductioncanbefoundin[8].
rierbeforetheupdateofbottom,guaranteeingthatthepushed Axiomaticexecutionwitnessescaptureabstractmemoryevents
elementcannotbestolenbeforebottomisupdated. associated with memory-related instructions and internal transi-
• Takentasksarereservedfirstbyupdatingbottom;again,inour tions of the model. Unlike in stronger models such as x86, each
ARMv7code,thesyncbarrierplacedaftertheupdatetobottom memoryaccessisrepresentedatrun-timebytwodistinctevents:an
willensurethatitwillnotbeconcurrentlystolen. issuingevent—calledsatforreadsandiniforwrites—eventually
• Stolen tasks are reserved by updating top. The only situation followedbyacommiteventwhenthespeculativestateofthein-
where steal and take contend for the same task is when the structionisresolved.Onceawriteinstructioniscommitted,events
deque has a single element left; this particular conflict is re- thatpropagateittootherthreadscanbeobserved—propagationto
solvedthroughtheCASinstructionsinbothtakeandsteal.This threadAisdenotedpp .Alltherelationspartofanexecutionwit-
A
scenarioallowedChaseandLevtomaketheCASintakecon- nessarelistedinTable1.
ditionaluponthesizeofthedequebeing1.Thecorrectnessof The core of the axiomatic model builds on the evord relation,
this optimization on a relaxed memory model depends on the modelingthehappens-beforeorderbetweenevents.Thissatisfies
presenceofthetwofullbarriersintakeandsteal,toensurethat thefundamentalproperty:
atleastoneoftheparticipantswillhaveaconsistentviewofthe
sizeofthedeque.Havingjustonetakeorstealseeingaconsis- −ev−o→rd ⊃ −a−ft→er ∪ −b−ef−o→re ∪ −c−om−→m ∪ −in→sn ∪ −lo−c→al
tentviewofthesizeofthedequeisenough:ifitistake,thatwill andmustbeacyclicforanexecutiontobeconsistent.
forceaCAStobeperformed;ifitissteal,theindexreservation We assume that the atomic sections, used to represent CAS-
willensureanemptyreturnvalue. like behaviors, are executed atomically and obey a total order.
• Finally, stolen tasks are protected from being concurrently We model them either as a single instance of a read instruction
stolen multiple times by the monotonic CAS update to top (failedCAS)oranatomicread–writepairofinstructioninstances
insteal.ThisCASordersstealoperationsandmakesthemmu- (successfulCAS).Theatomicityoftheseaccessesiscapturedby
tually exclusive. At the same time, steal operations that abort the −p−o-−at−o→m relation.Wedonotassumeanyotherpropertyonthese
duetoafailedCASdonotchangethestateofthedeque. atomic sections (e.g., cumulativity). In practice, atomic sections
canbeimplementedwithLL/SCinstructions.
2.2 ComparisonoftheC11andARMimplementations
Weuseseveralnotationshortcuts.Werefertothedequeglobal
OurC11implementationinFigure1isoptimalinthesensethatno variables top, bottom, and array as t, b, and a. Elements of the
C11 synchronization can be removed without breaking the algo- bufferarewrittenx ,whereiisthevirtualindexinnaturalnumbers
i
int take(Deque *q) { int take(Deque *q) {
size_t b = load_explicit(&q->bottom, relaxed) - 1; size_t b = R(q->bottom) - 1; (a)
Array *a = load_explicit(&q->array, relaxed); Array *a = R(q->array); (b)
store_explicit(&q->bottom, b, relaxed); W(q->bottom, b); (c)
thread_fence(seq_cst); sync;
size_t t = load_explicit(&q->top, relaxed); size_t t = R(q->top); (d)
int x; int x;
if (t <= b) { if (t <= b) {
/* Non-empty queue. */ x = R(a->buffer[b % a->size]); (e)
x = load_explicit(&a->buffer[b % a->size], relaxed); if (t == b) {
if (t == b) { bool success = false;
/* Single last element in queue. */ atomic /* Implemented with LL/SC. */
if (!compare_exchange_strong_explicit(&q->top, &t, t + 1, seq_cst, relaxed)) if (success = (R(q->top) == t)) (f)
/* Failed race. */ W(q->top, t + 1); (g)
x = EMPTY; if (!success) x = EMPTY;
store_explicit(&q->bottom, b + 1, relaxed); W(q->bottom, b + 1); (h)
} }
} else { /* Empty queue. */ } else {
x = EMPTY; x = EMPTY;
store_explicit(&q->bottom, b + 1, relaxed); W(q->bottom, b + 1); (i)
} }
return x; return x;
} }
void push(Deque *q, int x) { void push(Deque *q, int x) {
size_t b = load_explicit(&q->bottom, relaxed); size_t b = R(q->bottom); (a)
size_t t = load_explicit(&q->top, acquire); size_t t = R(q->top); (b)
Array *a = load_explicit(&q->array, relaxed); Array *a = R(q->array); (c)
if (b - t > a->size - 1) { /* Full queue. */ if (b - t > a->size - 1) { /* Full queue. */
resize(q); resize(q);
a = load_explicit(&q->array, relaxed); a = R(q->array); (d)
} }
store_explicit(&a->buffer[b % a->size], x, relaxed); W(a->buffer[b % a->size], x); (e)
thread_fence(release); sync;
store_explicit(&q->bottom, b + 1, relaxed); W(q->bottom, b + 1); (f)
} }
int steal(Deque *q) { int steal(Deque *q) {
size_t t = load_explicit(&q->top, acquire); size_t t = R(q->top); (a)
thread_fence(seq_cst); sync;
size_t b = load_explicit(&q->bottom, acquire); size_t b = R(q->bottom); (b)
int x = EMPTY; int x = EMPTY;
if (t < b) { if (t < b) {
/* Non-empty queue. */ Array *a = R(q->array); (c)
Array *a = load_explicit(&q->array, consume); x = R(a->buffer[t % a->size]); (d)
x = load_explicit(&a->buffer[t % a->size], relaxed); ctrl_isync;
if (!compare_exchange_strong_explicit(&q->top, &t, t + 1, seq_cst, relaxed)) bool success = false;
/* Failed race. */ atomic /* Implemented with LL/SC. */
return ABORT; if (success = (R(q->top) == t)) (e)
} W(q->top, t + 1); (f)
return x; if (!success) return ABORT;
} }
return x;
Figure1. C11codeofChase–Levdeque,withlow-levelatomics }
Figure2. ARMv7pseudo-codeofChase–Levdeque
beforeanywrap-aroundisapplied.Barrierinstructionsareomitted tecture.Tothatend,wepinpointspecificsubgraphsofanexecution
for brevity when implied by the presence of a −s−y→nc or −c−trl−-is−y→nc witness(hereafter,executiongraphs)thatcannotoccurtogetherin
relation.Irrelevantvaluesinreadsandwritesarereplacedwiththe the same consistent execution witness. We then show that all in-
placeholder“_”(e.g.,Rx,_).Wedonotlabelinstructioninstances correctexecutions,suchasthosecontainingtwoinstancesofsteal
individually, but decorate them with a disambiguating execution readingthesamevalueaddedbyasingleinstanceofpush,cannot
prefix, identified by adot. These prefixes do not only distinguish haveconsistentexecutionwitnessesand,assuch,cannothappen.
between instruction instances, but also group related instruction Theproofisstructuredasfollows.In4.1weprovidebasictech-
instances within a same execution unit (usually an invocation of nical definitions and properties of the memory model, which are
oneofpush,takeorsteal).Forthis,whennoprefixisspecified,the usedthroughouttheproof.In4.2wedescribeallthepossibleexe-
lastprefixinleft-to-rightorderisassumed. cutiongraphsforeachofthethreeoperations(push,takeandsteal),
followingthecontrolflowoftheARMv7codeinFigure2.In4.3
we show how the succession of dynamic arrays built by resizing
4. ProofofcorrectnessoftheARMv7code canbeabstractedasasinglesequenceofuniqueabstractvaluesin-
dependentofresizeoperations,withstrongcoherenceandconsis-
Theproofisdividedintofiveparts;itvalidatesthecriteria2to4
tencyproperties.Corollary2establishesCriterion2(well-defined
enumeratedinSection2.1.Sincepushandtakeneverexecutecon-
reads).In4.4webuildonthepreviousabstractiontoproveTheo-
currentlyandbisonlyevermodifiedinoneofthesefunctions,the
rem1,pertainingtotheuniquenessofelementstakenandstolen,
proofofCriterion1doesnotinvolvereasoningaboutconcurrency
whichcorrespondstoCriterion3(uniqueness).Finally,in4.5,we
andweomitithere.
relyonallpreviousresultstoproveTheorem2establishingCrite-
Theproofbuildsonapreciseanalysisofallthepossibleexe-
rion 4 (existence): the existence of matching take or steal opera-
cutionwitnessesofarbitraryinvocationsofthealgorithm.Were-
tionsforeverypushedelement,undertheappropriatehypotheses.
callthatanexecutionwitness,asdefinedbytheARMv7axiomatic
model,isagraphcapturingallmemoryeventsoccuringduringan
4.1 Preliminaryproperties
execution(vertices),aswellastherelationsthatlinkthem(edges).
Individuallemmasstrivetonarrowdownthesetofpossibleexecu- Before delving into the details of the proof itself, we introduce
tionwitnesses,basedonpropertiesofthealgorithmandthearchi- somesupportdefinitionsandrelatedproperties.
Rl,α readofvalueαfromlocationl(_standsforanyvalue) implies −p−p-−s→at , by definition of communication edges (if threads
Wl,α writeofvalueαtolocationl(_standsforanyvalue) aredifferent)oruniprocessorconstraints(ifsamethread).
sync memorybarrier(usuallyimpliedby −s−y→nc )
isync instructionbarrier(usuallyimpliedby −c−tr−l-−isy−→nc ) Lemma2. Thefollowingpropertiesinvolve −p−p-−s→at and −p−o-−l→oc :
sat(X) satisfy(a.k.a.complete)eventofareadinstruction (i) A.Wx,_−→rf B.Rx,_−p−o-−l→oc B(cid:48).Rx,_ =⇒ A.Wx,_−p−p-−s→at B(cid:48).Rx,_
ini(X) initializeeventofawriteinstruction (ii) A.Wx,_−→co B.Wx,_−p−p-−s→at C.Rx,_ =⇒ A.Wx,_(cid:54)−→rf C.Rx,_
com(X) commiteventofanin-flightorspeculativeinstruction (iii) Wx,_−p−p-−s→at Ry,_−→dp Rz,_ =⇒ Wx,_−p−p-−s→at Rz,_
ppA(X) propagatetothreadofAevent (iv) ¬(cid:0)A.Wx,_−p−p-−s→at B.Ry(cid:48),_−→dp B.Wx(cid:48),_−p−p-−s→at A.Ry,_−→dp A.Wx,_(cid:1)
−p→o programorder Proof. Weproveeachpointseparately:
−p−o-−at−o→m atomicoperationinprogramorder(forCAS;seebelow) (i) Ifthewriteandthereadshappeninthesamethread,thenallin-
−p−o-−lo→c same-locationaccessinprogramorder(definedin4.1) struction instances belong to that thread, and program order prevails.
−→co writecoherence Otherwise, either A.Wx,_−→rf B(cid:48).Rx,_ and the result is immediate, or
−→rf readfrom A.Wx,_(cid:54)−→rf B(cid:48).Rx,_ and B.Rx,_−p−o-−lo→c B(cid:48).Rx,_ implies the following:
−r→(cid:27) readfromfar(definedin4.3) com(B.Ry,_)−lo−c→al sat(B(cid:48).Ry,_),bydefinitionof −lo−c→al .Hence:
−→fr fromread
−a−d→dr addressdependence(usuallyimplicit) ppB(A.Wx,_)→sat(B.Ry,_)−in−→sn com(Ry,_)−lo−c→al sat(B(cid:48).Ry,_)
−c→trl controldependence(usuallyimplicit) (ii) Suppose A.Wx,_(cid:54)−→rf C.Rx,_. Then C.Rx,_−→fr B.Wx,_, and we
−d−a→ta datadependence(usuallyimplicit) havethefollowingcycleintheeventhappens-beforeorder:
−c−tr−l−d-→−ipsy−→nc onbosne-rcvuambuleladtievpeelnodceanlcoerd(deerfiinngedbainrr4ie.1r)(seebelow) sat(C.Rx,_)−c−om−→m ppZ(B.Wx,_)→sat(C.Rx,_)
−s−y→nc cumulativefullbarrier(seebelow) (iii) FollowsfromLemma1.
−p−p-−s→at write-to-readpropagation(definedin4.1) (iv) Assumethat:
−b−a−e−ff−to→e→rre abfetfeorrebabrarrireireredegdege A.Wx,_−p−p-−s→at B.Ry(cid:48),_−d→p B.Wx(cid:48),_−p−p-−s→at A.Ry,_−d→p A.Wx,_
−c−om−→m communicationedge IfA∼Bthenthereisacyclein −p→o .Otherwise,byLemma1,wehavea
−in−→sn intra-instructionorderedge cycleintheeventhappens-beforeorder:
−−elov−−oc→→ardl leovceanltohradpepreendsg-beeforeorder(usuallytypesetas→) ppB(Wx,_)→sat(Ry(cid:48),_)→com(Wx(cid:48),_)−in−→sn ppA(Wx(cid:48),_)
OnARMv7, −s−y→nc correspondstoadmbinstructionwhile −c−tr−l-−isy−→nc corre- →sat(Ry,_)→com(Wx,_)−in−→sn ppB(Wx,_)
Lemma3. Thefollowingpropertiesinvolvingbarriersapply:
spondstoadependentconditionalbranchfollowedbyanisbinstruction.
Table 1. Summary of relations used in the ARMv7 axiomatic (i) (Wx,_−s−y→nc Wy,_−p−p-−s→at Rz,_ ∨ Wx,_−p−p-−s→at Ry,_−s−y→nc Rz,_)
model =⇒ Wx,_−p−p-−s→at Rz,_
(ii) A.Wx,_−→rf B.Rx,_−s−y→nc B.Wy,_−p−p-−s→at C.Rx,_
=⇒ A.Wx,_−p−p-−s→at C.Rx,_
For convenience, we define the −p−o-−l→oc relation, which relates (iii) LetXstandforA.Wx,_−→rf B.Rx,_or(A∼B).Wx,_
local(same-thread)accessestothesamememorylocation; −p−o-−l→oc andY standforC.Wy,_−→rf D.Ry,_or(C ∼D).Wy,_
impliesan instruction-levelcommunicationedge −→co , −→rf or −→fr . thenthefollowingholds:
Inparticular, −p−o-−l→oc implies −→co betweentwowrites. ¬(X−s−y→nc B.Ry,_−→fr C.Wy,_∧Y −s−y→nc D.Rx,_−→fr A.Wx,_)
Wedefinethedependencerelation −→dp asfollows:
Proof. Weproveeachpointseparately:
Rx,_−→dp Ry,_ ⇐de⇒f Rx,_(−a−d→dr ∪ −c−trl−-is−y→nc )Ry,_ (i) If Wx,_ and Rz,_ occur in the same thread, then all instruction
Rx,_−→dp Wy,_ ⇐de⇒f Rx,_(−a−d→dr ∪ −c→trl ∪ −d−a→ta )Wy,_ instances belong to that thread and program order prevails. Otherwise,
supposeRz,_executesinA;wehavetwocases:
Lemma1.RTxh,e_f−→doplloRwyi,n_g=p⇒ropesrattie(Rsixn,v_o)l→vingsa−t→d(pRya,p_p)ly: ppA(Wx,_)−b−ef−o→re ppA(sync)−b−ef−o→re ppA(Wy,_)→sat(Rz,_)
Rx,_−→dp Wy,_ =⇒ sat(Rx,_)→com(Wy,_) Ortheotherwayaround:
Proof. In the case the of an address or control dependence, the result is ppA(Wx,_)→sat(Ry,_)−in−→sn com(Ry,_)−lo−c→al com(sync)−lo−c→al sat(Rz,_)
animmediateconsequenceofthedefinitionofintra-instructionandlocal Inbothcases,ppA(Wx,_)→sat(Rz,_).
orders.Itremainstobeshownthattheresultholdsfor −c−tr−l-−isy−→nc :adepen- (ii) Suppose A ∼ C. If A ∼ B, then program order prevails:
dentconditionalbranchinstruction,ctrl,followedbyanisyncbarrier.Sup- all the instruction instances belong to the same thread. If not, suppose
poseRx,_−c−tr−l-−isy−→nc Ry,_.Thenwehave:sat(Rx,_)−in−→sn com(Rx,_)−lo−c→al C.Rx,_−p→o A.Wx,_;thentheeventhappens-beforeordercontainsthefol-
com(ctrl)−lo−c→al com(isync)−lo−c→al sat(Ry,_). lowingcycle:
We define the relation −p−p-−s→at between instruction instances, ppB(A.Wx,_)−c−om−→m sat(B.Rx,_)−in−→sn com(Rx,_)−lo−c→al com(sync)
A.Wx,_−p−p-−s→at B.Ry,_,asfollows:1 −lo−c→al com(B.Wy,_)−in−→sn ppC(Wy,_)→sat(C.Rx,_)−in−→sn com(Rx,_)
(cid:40)Wx,_−→po Ry,_ ifA∼B −lo−c→al com(A.Wx,_)−in−→sn ppB(Wx,_)
ppB(Wx,_)→sat(Ry,_) ifA(cid:54)∼B Otherwise,supposeA(cid:54)∼C.IfA∼B,thenA.Wx,_−s−y→nc B.Wy,_andwe
havetheresultfrom(i).Ifnot,wehave:
wprheefirxeesAA∼andBBmbeelaonnsgtthoatthiensstarmucetitohnreainds.tances grouped under ppB(A.Wx,_)−c−om−→m sat(B.Rx,_)−in−→sn com(Rx,_)−lo−c→al com(sync)
Intuitively, −p−p-−s→at represents a “known-to” relation in the fol- Thus,wehaveppC(A.Wx,_)−b−ef−o→re ppC(sync)−b−ef−o→re ppC(B.Wy,_)→
lowingsense:A.Wx,_−p−p-−s→at B.Ry,_meansthat,atthetimeofread- sat((Ciii.)RSxu,_p)p.osethecontrary.IfB∼D,then −→rf and −→fr formapaththat
ingy,thatspecificwritetox(aswellasanywritethatiscoherence- goesagainst −p→o :thegraphisinvalidaccordingtouniprocessorconstraints.
beforeit)isknowntothethreadexecutingB.Itisclearthat −→rf Otherwise,B(cid:54)∼Dandthefollowingholds(omittingintermediatesteps
inelaborating −b−ef−o→re forconciseness):
1Notethat −p−p-−s→at doesnotimplyaneventhappens-beforeorder onthe • com(B.sync)−lo−c→al com(C.Wy,_)−lo−c→al com(D.sync)−in−→sn ppB(sync)
eventsmakinguptherelatedinstructioninstances. ifB∼C.
• com(B.sync)−lo−c→al sat(Ry,_)−c−om−→m ppB(C.Wy,_)−b−ef−o→re ppB(D.sync) relation. Initially β0 = 0. Since all push and take operations
otherwise. occurinasinglethread,andstealoperationsneveralterthevalue
Eafittehreredwgaey,bceotwme(eBn.tshyenct)wo→baprrpieBrs(:Dp.psyDn(cB)..sByyncd)e−fia−fnt→eitriocno,mw(eDh.asvyenca)n. oofrdbe,rtwheitheilnemtheenptsusohfa(nβdnt)akceororpesepraotniodntso.Swimritielasrltyo,bweindepfirongertahme
Moreover,eitherA∼DorA(cid:54)∼D:
sequence(τ )ofvaluestakenbythevariablet.Weassumeτ =0.
• ppD(B.sync)−a−ft→er com(D.sync)−lo−c→al com(A.Wx,_)−in−→sn ppB(Wx,_) Furthermorem,sinceallwritestotarefromCASinstructions,0which
ifA∼D.
• optphDer(wBis.es.ync)−a−ft→er com(D.sync)−lo−c→al sat(Rx,_)−c−om−→m ppD(A.Wx,_) atrbeysoeqnue,en(τtimal)lyisomrdoenreodto,naincdalalyllisnuccrehaCsiAngS,ainnsdtrsu.ct.tiτomns=inmcre.ment
Thus,inallcases,wehaveacycle: Foreachindexi,wedefinethesequence(ξiv)v∈Nofsuccessive
−c−om−→m satc(Bom.R(Bx,._s)y−inn−→csn)−bc−eof−mo→re(Rpxp,B_)(A−lo−c.→Wal xco,_m)(B.sync) uvWanlxdueie,rs_lygoiinfvgeanaprtrouastyhh.eOoepnleelrymatteihnoetn,alatrsietngdsauerxdchlieswisnrtiohtfeetidhseecqaauldeldebdryesstsihge&nilxfiasctaownftrtihatees
4.2 Executionpaths it induces a new value in an (ξiv) sequence, while writes due to
resizingdonot.Foralli,ξ0,thevaluebeforethefirstsignificant
We consider the three operations of the work-stealing algorithm: write to x location, is undiefined: ξ0 = ⊥. Similarly, a read is
take, push and steal. Each of them exhibits different execution i i
significantifitoccursinasuccessfulinstanceoftakeorsteal.
paths depending on control flow. Data and address dependences
are implicit in the notations and are omitted for brevity. Control Lemma4. Foralli,(ξv)isgloballycoherent.
i
dependencesareimpliedbytheguardconditionsineachcaseand
avraeriaablsleosocmarirttyeidn,gbtuhtewcoenetrxopllidceitpethnedecnocnes.trGairneteskolnetttehresbβ,anτd, ξt oPfrothoef.aGddivreensstwofotshiegnuinfidcearnltywinrgitaersrWay)x.iI,f_WanxdiW,_xa(cid:48)in,d_aWtixnd(cid:48)i,e_xbio(trhegwarridtelestos
thesamememorylocation,thentheyareorderedbywritecoherence.Ifthey
denotethememoryvaluesofb,t,andsomexi,respectively.Reads donot,thentheremustbearesizeoperationafterthefirstwriteandbefore
andwritesareannotatedwiththecorrespondinglinefromFigure2. thesecond(allwriteshappeninthesamethread).Becauseofthecumulative
For take and steal, we say that an instance of the operation barrierafteraresizeoperation,threadsthatseethesecondvaluemusthave
is successful if it returns one element; otherwise (including if it seenthefirstbeforehand.Hence,thereisaglobalcoherenceorderonthe
returnsempty)itisconsideredfailed. writes,whichcorrespondstotheorderofpushoperations.
Wedefinetherelationreadfromfarasfollows:forsomemem-
4.2.1 Take orylocationsm0,...,mn andsomevaluev,Wm0,v−→r(cid:27) Rmn,v
Twofailurecasesreturnnoelement(empty),andtwosuccesscases ifWm0,v−→rf Rmn,vorthereexistsasequenceofcopiescarrying
returnoneelementfromthedeque.Allfourpathsstartwith: thevalueofthewritetotheread:
(a)Rb,β−→po (b)Ra,&x−→po (c)Wb,β−1−s−y→nc (d)Rt,τ Wm0,v−→rf Rm0,v−d−a→ta Wm1,v−→rf ··· −d−a→ta Wmn,v−→rf Rmn,v.
Forconciseness,wehereafteromitthevariablenamefromreads
Specificcontinuationsforeachpatharelistedbelow. and writes whenever the variable can be inferred from the value:
ReturnemptywithoutCAS, β−τ ≤0: ··· −→po (i)Wb,β e.g., Wβn stands for Wb,βn. Let Wξiv denote the vth significant
Returnemptywith(failed)CAS, β−τ =1,τ (cid:54)=τ(cid:48): writeatindexi,andRξivasignificantreads.t.Wξiv−→r(cid:27) Rξiv.
··· −→po (e)Rxβ−1,ξ−→po (f)Rt,τ(cid:48)−→po (h)Wb,τ +1 Lemma5. GivenawriteWxi,_andareadRx(cid:48)j,_,
RReettuurrnnoonneewwiitthho(usutcCcAesSsf,uβl)−CAτS>, β1:−τ··=· −→p1o: (e)Rxβ−1,ξ i(cid:54)=j =⇒ Wxi,_(cid:54)−→rf Rx(cid:48)j,_
··· −→po (e)Rxβ−1,ξ−→po (f)Rt,τ−p−o-−at−o→m (g)Wt,τ+1−→po (h)Wb,β Proof. If the addresses of the underlying arrays differ, then the memory
locationsreadandwrittenaredistinctandtherecanbenoreadfromrelation.
4.2.2 Push
Otherwise,sinceoldarraysareneverreused,theaddressesarethesame
There are two paths: a straight case, and a resizing case which andi ≡ j mod size(x)Rx(cid:48)j,_belongstoasuccessfulinstanceoftake,
growstheunderlyingcircularbuffer. push(withresizing),orsteal.LetXbethatinstance.
LetP betheinstanceofpushtowhichWxi,_belongs.InP,wehave
Straight, β−τ <size(x)−1: thefollowingexecutiongraph:
(aR)eRsbiz,iβng−→p,oβ(b−)Rτt,≥τ−s→pioze((cx)R)a−,&1:xw−→phoer(ee)xW(cid:48)rxeβfe,rξs−ts−oy→ncth(efn)eWwba,rβra+y 1 P.Rt,τP−c→trl Wxi,_−s−y→nc Wb,βP +1
(a)Rb,β−→po (b)Rt,τ−→po (c)Ra,&x−→po resize where τP ≤i≤βP and βP −τP <size(x)−1
−s−y→nc (d)Ra,&x(cid:48)−→po (e)Wx(cid:48)β,ξ−s−y→nc (f)Wb,β+1 Letusassumei(cid:54)=j∧Wxi,_−→rf Rx(cid:48)j,_andshowitisindeedimpossible.
where resize =Rxτ,ξτ−→po Wx(cid:48)τ,ξτ−→po ··· Assume X is a successful instance of take or push. Since X and
−→po Rxβ−1,ξβ−1−→po Wx(cid:48)β−1,ξβ−1−s−y→nc Wa,&x(cid:48) dPerbe(tlhoengortdoertheofsalomaedstharneadd,stPoremsutosttohcecusrambeefolroecaXtioninispropgrersaemrveodr-:
4.2.3 Steal P.Wxi,_−p−o-−lo→c X.Rx(cid:48)j,_).
Ifj<i,thenj≤i−size(x).However,thefollowingmustholdinP:
There are three paths: two failure cases and one success case.
Failurereturnsnoelementandsuccessreturnsastolenelement. τP ≤i≤βP ∧βP −τP <size(x)−1
ReturnemptywithoutCAS, β−τ ≤0: (a)Rt,τ−s−y→nc (b)Rb,β hence j<i−size(x)+1≤βP −size(x)+1<τP
Returnemptywith(failed)CAS, β−τ >0∧τ (cid:54)=τ(cid:48): Furthermore,ifX isatakeoperation,Rx(cid:48)j,_readsthelastelementof
(a)Rt,τ−s−y→nc (b)Rb,β−c−trl−-is−y→nc (c)Ra,&x−→po (d)Rxτ,ξ−c−trl−-is−y→nc (e)Rt,τ(cid:48) the deque, and j = βX −1 ≥ τX; if X is a push operation, Rx(cid:48)j,_
Returnonewith(successful)CAS, β−τ >0: resultsfromacopyoperationoftheresizingcode,hencej ≥ τX.Since
(a)Rt,τ−s−y→nc (b)Rb,β−c−tr−lc−-−tisr−yl−-→nis−cy→n(cc)(Re)aR,t&,τx−−p→p−oo-−a(t−od→m)R(xfτ)W,ξt,τ +1 XP.Rotc,cτuPrs−p−aof-−lto→ecrXP.Rint,pτrXogaranmdjo<rdeτrPan≤dτtXis≤mjo.nIomtopnoiscsaiblllye.increasing,
Ifi < j,then,sincej ≥ βX,bmustincreasefromβP +1toj+1
4.3 Significantreadsandwrites betweenthewriteinP andthereadinX.Hence,theremustbeaninstance
P(cid:48)ofpushbetweenPandX(inprogramorder)thatincrementsbtoj+1.
We define the sequence (βn) of values taken by the variable b Indeed,theonlywritesthatincreasethevalueofboccurinpushandtake;
overthecourseoftheprogram,accordingtothewritecoherence andtheeffectoftakeasawholeneverincreasesthevalueofbsinceitfirst
decrementsthevariable.Wehave: Otherwise, 0 < v < u. Let W.Wξv be the significant write s.t.
i
henPc.eWxPi,.W_−p−xo-−ilo→,_c−→cPo(cid:48)P.W(cid:48).Wxjx,_j−p,−_o-−−plo−→pc-−s→aXt.XRx.R(cid:48)jx,_(cid:48)j,_ WcWar.(cid:48)rW.yWiξnivxgi−r→,t(cid:27)hξeivRv−→xrfail,uRξexivio.,fIξniξviv.oMtthooerrReowxvoie,rrdξ,siav,c.ctTohrehdraeitngesextoqisutthesneacdeesfieeqnnuidteisonncweoitfho(fξaivco)wparinietdes
Thus,fromLemma2(ii),P.Wxi,_(cid:54)−→rf X.Rx(cid:48)j,_. Wtheξuseimnapnrtoigcsraomforredseizr.ing, W.Wξiv and W(cid:48).Wxi,ξiv must come before
execNuotiwo,nagsrsaupmhefoXrXis:asuccessfulinstanceofsteal.Wehavethefollowing iWehavetwocases:eitherWξiuandRxi,ξivrefertothesamememory
locationortheydonot.
X.Rt,τX =j−s−y→nc Rb,βX−c−tr−l-−isy−→nc Ra,&x(cid:48)−p→o Rx(cid:48)j,_ Assumethattheyrefertothesamememorylocationxi.Thenitmust
−c−tr−l-−isy−→nc Rt,τX−p−o-−at−o→m Wt,τX+1 bethatW(cid:48).Wxi,ξiv−p−o-−lo→c Wxi,ξiu,andwehave:
Ifj<i,thenj≤i−size(x).However,thefollowingmustholdinP: W(cid:48).Wxi,ξiv−→co Wξiu−p−p-−s→at Ra,&x−a−d→dr Rxi,ξiv
j<i−size(x)+1≤βP −size(x)+1<τP HenCceo,nfvreormselLye,masmsuam2e(tihi)a,tWthe(cid:48).yWdoxin,oξtivre(cid:54)−→rfferRtoxit,hξeivs.aImmepmosesmibolery. location.
HenceτX =j<τP.Sincetincreasesmonotonically,itmustbethat: ThentheremustbearesizeoperationbetweenW(cid:48).Wxi,ξivandWξiu:
X.Rx(cid:48)j,_−c−tr−l-−isy−→nc Rt,τX−p−o-−at−o→m Wt,τX+1 Wa,&x−s−y→nc W(cid:48).Wxi,ξiv−s−y→nc Wa,&x(cid:48)−s−y→nc Wx(cid:48)i,ξiu
−→rf Rt,_−s−y→nc Wt,_−→rf ···−s−y→nc Wt,τP−→rf P.Rt,τP−c→trl Wxi,_ −p−p-−s→at Ra,&x−a−d→dr Rxi,ξiv
HenceX.Rx(cid:48)j,_mustbecommittedbeforeWt,τX +1.SinceWt,τX + Hence,fromLemma3(i),Wa,&x−→co Wa,&x(cid:48)−p−p-−s→at Ra,&x.Andfrom
1 is (cumulatively) propagated to Wxi,_, X.Rx(cid:48)j,_ must be committed Lemma2(ii),Wa,&x(cid:54)−→rf Ra,&x.SincethereisonlyonewriteWa,&x
1be−pf−po-−rse→atWPx.Ri,t_,.τFPo.rImfaWllyx:i,i_t−→rffolRloxw(cid:48)js,_ftrhoemnWLexmim,_a−p−p3-−s→a(tii)Rtxh(cid:48)ja,t_.WWte,τgXet:+ tohfatc(ogiipi)vieeTsshtcehareerrveyaxilniusgets&thaxewtvroiatlaeu,eWwoe.Wfhξaξvvivetsao.tc.RoWnξtv.rWa.dTξichivtai−otr→(cid:27)ns.eRqξuiven,caendenadsseqwuietnhcae
X.Wt,τX+1−p−p-−s→at P.Rt,τP−c→trl Wxi,_ writeW(cid:48).Wξv−→rf Rξv.Sinceu ≤i v,Wξiu−p→o W.Wξv bydefinitionof
i i i i
∧P.Wxi,_−p−p-−s→at X.Rx(cid:48)j,_−c−tr−l-−isy−→nc Wt,τX+1 F(ξriovm).TLheamnmksat3ot(hi)e,bwaerrgieertaWfteξruW−p−pξ-−ius→atinRpξuvsh.,Wξiu−s−y→nc W(cid:48).Wξiv−→rf Rξiv.
Lemma2(iv)tellsthatitisimpossible.ThusP.Wxi,_(cid:54)−→rf X.Rx(cid:48)j,_. i i
If i < j, then i ≤ j −size(x), and there must be an instance P(cid:48) Corollary 2 (Well-defined significant reads). Given a significant
ofpushs.t.P(cid:48).Wb,j+1−p−o-−lo→c Wb,βX−→rf X.Rb,βX (sothatindexjbe readRxi,ξ,ξ=ξivforsomev>0.
accessibleinX).P(cid:48)cannotoccurbeforeP inprogramorderbecause,as
above,wewouldhaveτP(cid:48) ≤ τP ≤ iontheonehand,andi ≤ j − ProoSfu.pLpeotseXξb(cid:54)=etξhve,stuhcecnesξsf=ul⊥insctaanncoenolyftbaekeanorusntdeaefilns.et.dRvxaliu,eξf∈romXt.he
sinizcere(axs)es<inτsPiz(cid:48)eo,nsothteheotihneerquhaalnidty.TsthilelhuonlddesrliyfitnhgeasrirzaeysaolfsPomanodnoPto(cid:48)ndicifaflelry. uninitializedarray,ipriortocopying.Indeed,ifxiisnotaffectedbycopying,
thenitmustbeoneofthenewslotsallocatedbytheresizing,henceitsinitial
HenceP(cid:48)occursafterP.FurthermoreWx(cid:48)j(cid:48),_∈P(cid:48).IfxinP andx(cid:48)(cid:48)in valueisξ0.LetRbethepushoperationthatallocatesthearrayx.There
P(cid:48)refertodifferentarrays,thenaresizeoperationRmustprecedeP(cid:48),s.t. existsaξuisuchthat:
i
Wa,&x−p−o-−lo→c P.Ra,&x−p−o-−lo→c R.Wa,&x(cid:48)(cid:48) Wxi,⊥−→co R.Wxi,ξiu−s−y→nc Wa,&x−→rf X.Ra,&x−a−d→dr Rxi,ξ
−s−y→nc P(cid:48).Wx(cid:48)j(cid:48),_−s−y→nc Wb,j+1 It follows from Lemmas 2 (iii), 3 (i) and 2 (ii) that Wxi,⊥(cid:54)−→rf Rxi,ξ.
−p−o-−lo→c Wb,βX−→rf X.Rb,βX−c−tr−l-−isy−→nc Ra,&x(cid:48)−a−d→dr Rx(cid:48)j,_ ImpHosesnicbele,.ξ = ξv.WehaveRb,β ∈ X andβ ≥ i+1 > 0,forX is
hence Wa,&x−→co R.Wa,&x(cid:48)(cid:48)−s−y→nc Wb,βX−p−p-−s→at X.Rb,βX successful.Hence,ithereisaninstanceofpushP s.t.P.Wb,β−→rf X.Rb,β.
FromLemma2(iii),Wb,βX−p−p-−s→at X.Ra,&x(cid:48);Lemma2(ii)concludes Sinceβ ≥ i+1,eitherβ = i+1andWξiu ∈ P,ortheremustbe
that Wa,&x(cid:54)−→rf X.Ra,&x(cid:48). Since all resize operations allocate new ar- aninstanceofpushthatcontainsasignificantwriteWξiu andcomesbe-
rays,&x(cid:48) (cid:54)= &x,whichcontradictsourpremises.Otherwise,xandx(cid:48)(cid:48) foreP inprogramorder.Inbothcases,Wξiubelongstoapushoperation,
refertothesamearray,henceWxi,_−p−o-−lo→c Wx(cid:48)j(cid:48),_,andweget: phuenshce,Wuξ>u−s−0y→n.cMPo.rWeobv,eβr,.tIhfaXnkisstaontihnestbanarcreieorfataftkeer,aP.sWignbi,fiβca−p→notXwr.iRteξvin;
P.Wxi,_−p−o-−lo→c−→rfPX(cid:48).W.Rxb,(cid:48)j(cid:48)β,_X−s−y−c→n−tcr−l-W−isy−→nbc,jR+x(cid:48)j1,_−p−o-−lo→c Wb,βX oPhte.hnWecrebw,,ibβsye−i,pL−pe-P−sm→a.tWmXabs,.Rβ3ξ−→(rivif).aXInn.dRb6bo,,t0hβ<c−c−atrs−ule-−iss≤y,−→nWcv.Rξiuξiv−s−y→nacndP.WLebm,mβa−p−p3-−s→at(iiX).gRivξiievs,
ItfollowsfromLemmas3(i)and2(iii)that:
P.Wxi,_−→co Wx(cid:48)j(cid:48),_−p−p-−s→at Rx(cid:48)j,_ 4.4 Uniquenessofsignificantreads
Hence,fromLemma2(ii),Wxi,_(cid:54)−→rf Rx(cid:48)j,_. Trehaedsreastudltisffferroemnttihnedepxreesvicoaunsnsoetcrteiotrnieevsetathbelisshamtheateltewmoesnitgξnvifi.cTahnet
i
Corollary1. GivenasignificantwriteWξivandasignificantread onlypossiblecauseofduplicatesignificantreadsisthusreducedto
Rx(cid:48)j,_: i(cid:54)=j =⇒ Wξiv(cid:54)−→r(cid:27) Rx(cid:48)j,_. thecasewherethereadsaccessthesameindexi.
Proof. If i (cid:54)= j, we know that Wξv(cid:54)−→rf Rx(cid:48),_. Furthermore, all copies, Theorem1(Work-stealing:uniquenessofsignificantreads). Given
i j
whichhappenduringaresizeoperation,copyfromandtothesameindex. aworkerthreadexecutingasequenceof pushandtakeoperations,
Sincetherearelesscopiesthanthesizeoftheexpandedarray,therecanbe andfinitenumbernumberofthiefthreadseachexecutingstealop-
notwocopieswritingtothesamememorylocationinthenewarray.Hence, erations, all against a same deque. If X and Y are two distinct
therecanbenosequenceofcopiesfromWξvtoRx(cid:48),_.
i j successfulinstancesof stealortake,
Lemma6. GivenasignificantwriteWξiu andasignificantread ∀Rξv ∈X,∀Rξv(cid:48) ∈Y,i(cid:54)=i(cid:48)∨v(cid:54)=v(cid:48)
Rξv: i i(cid:48)
i
(i) Wξiu−p−p-−s→at Ra,&x−a−d→dr Rxi,ξiv =⇒ u≤v Lemma7. GivenS1andS2distinctsuccessfulinstancesof steal,
(ii) 0<u≤v =⇒ Wξiu−p−p-−s→at Rxi,ξiv ∀Rξiv ∈S1,∀Rξiv(cid:48)(cid:48) ∈S2,i(cid:54)=i(cid:48)
Proof. Weproveeachpointseparately: Proof. Allwritestotatomicallyincrementit(byatomicityofCAS).Hence
(i) Supposev<u.WedefineW(cid:48).Wxi,ξivasfollows. twosuccessfulstealoperationscannotwrite(thusread)thesamevalueof
Ifv =0,ξivisanundefinedvalue;letW(cid:48).Wxi,ξi0−→rf Rxi,ξivbethe t.Readsfromxinastealoperationaccesstheindexgivenbythevalueof
initializationofxi.W(cid:48).Wxi,ξi0comesbeforeWξiuinprogramorder. thetvariable.HenceRt,i∈S1andRt,i(cid:48)∈S2implyi(cid:54)=i(cid:48).
Lemma 8. Given T a successful instance of take and P an in- 4.5 Existenceofsignificantreads
stanceof push.IfP comesafterT inprogramorder,then:
Theorem 2 (Work-stealing: existence of significant reads). Con-
∀Rξv ∈T,∀Wξu ∈P,i(cid:54)=j∨v(cid:54)=u sideraworkerthreadexecutingasequenceof pushand takeop-
i j
erations,andafinitenumberofthiefthreadseachexecutingsteal
Proof. Assume i = j ∧ v = u. We have Rξv−p→o Wξu; therefore
i j operations,allagainstasamedeque.Ifthenumberof pushisfi-
Wξu(cid:54)−p−p-−s→at Rξv.FromLemma6(ii),itfollowsthatu > v.Wehavea
j i nite,thenallthreadsreachastationarystatewhereb=tinafinite
contradiction.
numberoftransitions,andthefollowingholdsglobally:
Lemma9. GivenT andT distinctsuccessfulinstancesof take,
1 2 ∀ξv,v>0 =⇒ ∃!Rξvinsomethreadbeforethestationarypoint
∀Rξv ∈T ,∀Rξv(cid:48) ∈T ,i(cid:54)=i(cid:48)∨v(cid:54)=v(cid:48) i i
i 1 i(cid:48) 2
Let P be the last instance of push in the worker thread, in
F
Proof. Wehavethefollowingexecutiongraphs: programorder.LetWβ ∈P andRτ ∈P .Wesaythatan
T1.Rβn−p→o Ra,_−p→o Wb,βn−1−s−y→nc Rt,τ−p→o Rξβvn−1−p→o ··· instanceXoftakeorstneFalistraFilingifRmβFn≥nFF∈X.
T2.Rβn(cid:48)−p→o Ra,_−p→o Wb,βn(cid:48)−1−s−y→nc Rt,τ(cid:48)−p→o Rξβvn(cid:48)(cid:48)−1−p→o ··· Lemma 11. Given X a successful trailing instance of take or
Andβn−1=iandβn(cid:48)−1=i(cid:48). steal: Rτm ∈X =⇒ m≥mF.
Sinceallinstancesoftakeoccurintheworkerthread,wehaveeither:
Proof. Wehavetwocases:
T1.Wb,βn−1−p−o-−lo→c T2.Rβn(cid:48) or T2.Wb,βn(cid:48)−1−p−o-−lo→c T1.Rβn • Assume X is an instance of take. X follows PF in program order:
Letusassumethefirstcaseaswellasi =i(cid:48)∧v =v(cid:48)andshowitis PF.RτmF −p−o-−lo→c X.Rτm,andm≥mF byuniprocessorconstraints.
• AssumeX isaninstanceofsteal.SinceX issuccessful,X contains
iβamnnpi(cid:48)nHo−ssetsani1nbc,celeae(,nPβtdhneTo)f1omp.tWhuuessrbht,citniahsc−aper−toe-wb−alo→esrcieintefTgrs2osW.myRmβbim,ktoie−p+ti−or+i-−cl1o→a1c.l.bTWe2twe.Rehbea,nvien+βa1nn,d−s.nt1.(cid:48)n;=th<eirk=ee≤xi(cid:48)isn=ts(cid:48) fmaosululsoctwcyeiinsesglfduretlhaiednssXtaam.nRceet,v_oa,flauanedC.DAthuSeebitnaosrtrtrihueecrtbbioaenrfro,ireherePnbceFetw.tWheeeβntnwXFo.,Rrwebae,d_hsaafnvrdoe:mthet
andβk−1 =iandβk =i+1(asnotedabove,takeasawholedoesnot WτmF −→rf PF.RτmF −s−y→nc WβnF −p−o-−lo→c Wβn−→rf X.Rβn−s−y→nc Rτm
increasethevalueofb).Wegetthefollowinggraph:
Rb,i−p→o P.Wξiu−s−y→nc Wb,βk=i+1−p−o-−lo→c T2.Rβn(cid:48)−p→o Ra,_−a−d→dr Rξiv(cid:48) FthreonmfoLlleomwmsafro3m(iiL),emwmeah3av(ei)Wthτamt FW−τp−pm-−sF→at−pX−p-−.s→RatβRnτ−sm−y→n.cTRoτtaml.orI-t
LLeemmmmaa63((i)i)thyaiteluds≤Pv.W(cid:48)aξniud−fpr−po-−ms→atLRema,m_−aa−d8→drthRatξviv(cid:48)<. Iut.thImenpofosslliobwles. from dmeFr ,oWn τCkA−→cSo iWnsτtrmucFtio∧nsWanτdk(cid:54)−→LrfemRmτma.2Th(eiir)efgourea,ramnte≥emthFat.∀k <
Lemma12. GivenX andY distinctsuccessfultrailinginstances
Lemma10. GivenT asuccessfulinstanceof takeandS asuc-
cessfulinstanceof steal, of takeorsteal,then: ∀Rξiv ∈X,∀Rξiv(cid:48)(cid:48) ∈Y,i(cid:54)=i(cid:48).
∀Rξv ∈T,∀Rξv(cid:48) ∈S,i(cid:54)=i(cid:48)∨v(cid:54)=v(cid:48) Proof. Assumei=i(cid:48).AccordingtoTheorem1,v(cid:54)=v(cid:48),hencethereexist
i i(cid:48) twodistinctsignificantwritesWξv andWξv(cid:48).Withoutlossofgenerality,
PTw−rs.i−yotRh→noβcfβ.nRnW−βp→−oen(cid:48)hR1a−ca−=vtr,−el_-−iit−sph→y−aoe→nncfdWoRτlmblao,,w(cid:48)β_i=n−pn→og−i(cid:48)Re.1xξe−τvscm−y(cid:48)u→n(cid:48)tci−oc−Rtnr−τl-g−misry−a→np−p→chosR:Rτmξβv(cid:48)n−p−−o-1−at−−po→→om ·W··t,−p→τomS(cid:48).+Rτ1m(cid:48) lwdeetrri.uteFs,uaarsntshdueWmbrmeeξfoiovv(cid:48)rr(cid:48)ee<−,s−Pyt→nhvFce(cid:48)r.;PeWPFisβF.anW.WFcβu.βmnSnFuiFnla−cpit−eooiv-−cXleo→ccubrraWesraradbifes,tir_e(cid:48)fri−→rrnofbmpoYuthP.sRhFwbar.,fWi_ttee−d→rsβp,eniaRFncξh,pivwrsoieggnrhaiafimvceao:nr-t
Letusassumei=i(cid:48)∧v=v(cid:48).Thenτm(cid:48) =i(cid:48) =i=βn−1.ForS HencewehaveWξv(cid:48)−p−p-−s→at Y.RξvfromLemma3(i)andLemma2(iii).It
tosucceed,wemusthaveτm(cid:48) <βn(cid:48).Hence,βn≤βn(cid:48). thenfollowsfromLie(cid:48)mma6thatv(cid:48)i≤v;thus,v<v(cid:48)≤v.Impossible.
Also,forT tosucceed,wemusthaveτm<βn.Twocases:
• If βn = τm + 1, then a successful CAS occurs in T. Moreover, Corollary3. Thecombinednumberofsuccessfultrailinginstances
βn =τm+1impliesτm(cid:48) +1=βn =τm+1,henceτm(cid:48) =τm. of takeandstealislessthanorequaltoβ −τ .
Impossible,sincetismonotonicallyincreasingandSmustalsocontain nF mF
asuccessfulCASwiththesamevalueoft. Proof. LetXbeasuccessfultrailinginstanceoftakeorsteal,andRβn ∈
• Itofnβinca>llyτimncr+ea1s,etsh,etnhenroeCmAuSstobcectuwrsoinwTriteasndAm.W(cid:48)τ>mm−→co.SBin.cWeτtmm(cid:48)osn.ot.- X(froamndLRemτmma∈11X)..HWeneckenτomw≥thaτtmnF≥. nF (bydefinition)andm ≥ mF
A.Wτm−→rf T.Rτm−→fr B.Wτm(cid:48)−→rf S.Rτm(cid:48)−s−y→nc Rβn(cid:48) Furthermore,atakeoperationalwayscontainsonedecrementingwrite
tob(byone),whichmaybefollowedbyoneincrementingwritetob(by
IfS.Rβn(cid:48)B−→fr.WT.τWm(cid:48)b−→,rfβnS.−Rτ1m,(cid:48)th−se−yn→ncwReβhnav(cid:48)e−→f:r T.Wb,βn−1 one)T.hHeerenfcoeren,≥XncFanimonpllyiesreβand≤atβannFin.dex i, s.t. τmF ≤ i < βnF.
∧T.Wb,βn−1−s−y→nc Rτm−→fr B.Wτm(cid:48) Lemma12tellstherecanbenomorethanβnF −τmF suchX.
ImpossibleaccordingtoLemma3(iii).ThereforeWβn(cid:48),thesource Lemma13. Thereisafinitenumberofsuccessful(trailingornon-
ofS.Rβn(cid:48) mustcomebeforeWβn+1 (incoherenceorder,hencein trailing)instancesof takeorsteal.
programorderasbothoccurinthesamethread).Consequently,(βn)
mustincreasefromβn−1=itoβn(cid:48)betweenn+1andn(cid:48).SinceT Proof. ItfollowsfromCorollary3thatthereisafinitenumberofsuccessful
doesnotincrementthevalueofb(executionwithoutCAS),theremust trailinginstancesoftakeorsteal.
beaninstanceP ofpushthatwritesP.Wβk−p→o Wβn(cid:48)−→rf S.Rβn(cid:48),s.t. Furthermore,theremustbeafinitenumberofnon-trailingtakeopera-
n<k≤n(cid:48)andβk−1=iandβk=i+1. tions,whichcomebeforePF inprogramorder.
Wegetthefollowingexecutiongraph: Lastly, there is a finite number of push operations, thus (βn) has a
P.Wξiu−s−y→nc Wb,i+1−p→o Wβn(cid:48)−→rf S.Rβn(cid:48)−c−tr−l-−isy−→nc Rξiv(cid:48) mvaaluxeimsuomf,tβlmesasxt.hSainncseomtweovsaulcuceesosffubl,sttheearleopcearnatbioennsommusotrreeathdadnifβfemreanxt
HencewehaveWξu−p−p-−s→at Rξv(cid:48)fromLemma3(i)andLemma2(iii). successfulinstancesofsteal.
i i Hencethefinitenumberofsuccessfulinstancesoftakeorsteal.
Finally,itfollowsfromLemma6thatu≤v(cid:48),andfromLemma8that
v<u≤v(cid:48).Wehaveacontradiction. Lemma14. Ineachthread,thereexistsXafailedinstanceof take
Theorem1followsdirectlyfromLemmas9,10and7. orsteals.t.∀Rβ ∈X,∀Rτ ∈X,β ≤τ .Furthermore,each
n m n m
threadmakesnomorethan1+m +β −τ attemptsattake straightforward fashion. In all cases that do not involve cumula-
orstealthatresultinafailedCASFinstruncFtion.2mF tivity, the −p−p-−s→at relation (defined in 4.1) combined with depen-
dences,whichformthecoreoftheARMv7/POWERproof,maybe
Proof. ItfollowsfromLemma13thatthereisafinitenumberofsuccessful
replacedwithanalogouspropertiespertainingtotheC11happens-
instances,henceafinitenumberperthread.Thus,theremustexistafailed
beforerelationcombinedwithrelease–acquiresemantics.Theone
instanceoftakeorsteal.
notabledifferencebetweenthetwomodelsliesintheabsenceofcu-
A failure can occur either because the deque is empty (βn ≤ τm)
or because of a failed CAS instruction. Suppose there is no X where mulativityinthedesignoftheC11abstractmachine:neitherC11
βn ≤ τm;thenallfailuresmustbeduetoafailedCASinstruction.A fences nor C11 atomic accesses guarantee cumulativity. A simi-
failedCASoccursifthetwovaluesoftreadduringtheinstanceXdiffer. lareffectcanbeachievedbychainingalternatingrelease–acquire
LetY1andY2betwosuchfailedinstancesexecutinginasamethread;let writesandreads,whichformahappens-beforepath.Butthisdevice
usassumethatY2followsY1inprogramorder,n1(cid:54)=n(cid:48)1andn2(cid:54)=n(cid:48)2: doesnotworkinsituationswherepropagationneedstobeasserted
Y1.Rτn1−p−o-−lo→c Rτn(cid:48)1−p−o-−lo→c Y2.Rτn2−p−o-−lo→c Rτn(cid:48)2 sbiettuwateieonntowcocurresaidns,threatshteeralthoapneraatrieoand.Ifnoflloorwmeadllyby(saeewLrietme.m3 Tah1i0s
TwheerheaveexisWtsτanw(cid:48)1r−pi−pte-−s→aWt τYn2(cid:48)1.R−→rτfnR2τ,na(cid:48)1nd−p,−o-−alo→scinY2th.Reτpnr2o.oDfuoeftLoeLmemmam1a12, (wi)e, ftaokrethdeofonromtraeladdes“corlidp”tivoanl)u,eitsmofubstobthebthoattttowmoacnodntcoupr,rewnhtesrteea“loalndd”
deducethatn(cid:48)1≤n2. couldbedefinedas“olderthanthevalueknowntotheotherparty
Sincen1(cid:54)=n(cid:48)1 ∧ n2(cid:54)=n(cid:48)2,andtismonotonicallyincreasing,itmust incoherenceorder”.Thepresenceofthetwocumulativebarriersin
bethatn1 <n(cid:48)1 ≤n2 <n(cid:48)2.HencesuccessiveCAS-failinginstancesin stealandtakeonARMv7guaranteesuchacondition:
asamethreadreadincreasingvaluesoft.ItfollowsfromCorollary3thatt
• ifthetakebarrierisorderedbeforethebarrierinsteal,thenthe
takesnomorethan1+mF +βnF −τmF differentvalues. program-order-previouswritetobottomwillbepropagatedto
Therefore,therecanbenomorethan1+mF+βnF−τmF CAS-failing theinstanceofsteal;
instancesoftakeorstealperthread.Sincethereisalsoafinitenumberof
successfulsuchinstances,anyfurthertakeorstealoperationsmustreturn • conversely, if the steal barrier is ordered before the barrier in
empty,andthethreadreachesitsstationarypoint. take,thenvaluereadbytheprogram-order-previousreadfrom
topwillbepropagatedtotheinstanceoftake.
Corollary4. Thecombinednumberofsuccessful(trailingornot)
In the second case, it is important to remark that the write that
instancesof takeandstealisequaltothenumberof push.
produced the value read in steal might belong to another thread,
Proof. Asuccessfulinstanceoftakeeitherdecreasesthevalueofbbyone and thus not be sequenced before the barrier. In the absence of
orincreasesthevalueoftbyone;asuccessfulinstanceofstealincreases cumulativity,itneednotbepropagatedtotheinstanceoftake.
thevalueoftbyone.Aninstanceofpushincreasesthevalueofbbyone. ToenforcethisparticularcaseofcumulativityinC11,werely
Itfollowsfromthepreviouslemmathattheworkerthreadreachesa on the properties of sequential consistency. By making all writes
stationarypointwhereb = t.Clearly,thiscannotoccurbeforeallpush (actually, CAS operations) to top sequentially consistent, we en-
operationsandallsuccessfulinstancesoftakehaveoccurred. surethatthereisatotalorderingbetweenthetwofences(intake
Sinceb = tatthestationarypointandallincreasestobprecede,the
andsteal)andthewritethatproducedthevalueoftopreadinthe
sumofincreasestotanddecreasestob(thecombinednumberofsuccessful
instanceofsteal.Furthermore,ifthatreadusesacquiresemantics,
instancesoftakeandsteal)mustbeatleastequaltothenumberofincreases
thenthereisahappens-beforerelationbetweenitandthestealbar-
tob(thenumberofpushoperations).Itisexactlyequal,asotherwisethere
wouldbemoresignificantreadsthansignificantwrites,whichisimpossible rier.Hence,thewritemustcomebeforesaidbarrierinsequential
accordingtoTheorem1. consistencytotalorder.Then,eitherthebarrierinstealisordered
beforethebarrierintake,ortheotherwayaround:
OnemayfinallyproveTheorem2.Ontheonehand,Corollary4
• if the steal barrier is ordered before the barrier in take, then
tellsthatthenumberofsignificantreads(fromasuccessfulinstance
itfollowsfromseq_cstbarriersemanticsthatthevalueoftop
oftakeorsteal)isequaltothenumberofsignificantwrites(from
readbytakecannotbeolderthantheonereadinsteal;4
an instance of push). On the other hand, Theorem 1 states that
• conversely,ifthetakebarrierisorderedbeforetheoneinsteal,
significantreadsuniquelymaptosignificantwrites.Byinjectivity,
thenthevalueofbottomreadbystealcannotbeolderthanthe
thereexistsauniquesignificantreadforeachsignificantwrite.
onewrittenintake.5
5. OntheC11implementation
6. Experimentalresults
The sequentially consistent implementation is a direct translation
Wepresentexperimentalresultsonthreecurrentandwidelyused
of the original algorithm using C11 seq_cst atomic variables for
architectures: (1) a Tegra 3 ARMv7 processor rev 9 (v7l) with 4
all shared accesses. It is obtained from the code in Figure 1 by
coresat1.3GHzand1GBofRAM;(2)anIntelCorei7-2720QM
replacingallmemoryorderconstantswithseq_cst;doingsomakes
machine with 4 cores (hyper-threading disabled) at 2.2GHz and
fencesunnecessary,hencetheyshouldberemoved.
4GBofRAM;and(3)adual-socketAMDOpteronMagny-Cours
TheoptimizedC11implementationimprovesupontheprevious
6164HEmachinewith2×12coresat1.7GHzand16GBofRAM.
version by replacing sequential consistency with release–acquire
All tests were compiled with GCC 4.7.0, the first release of
pairswhereappropriate.Itestablisheshappens-beforerelationsbe-
GCCtointroducebuilt-insupportforC11atomics.
tween reads and writes, as required by the proof. Unfortunately,
without relying on seq_cst, using only release, acquire and con-
3C11 defines a happens-before relation, which does not fully encapsu-
sumeoperations,wewereunabletoreproducetherequiredmemory
late the notion of cumulativity. The only inter-thread edges in happens-
orderingconstraintsneededonthePOWERandARMv7architec- beforecomefromwrite–readpairswithrelease–acquiresemantics(see[6]
tureswhileadheringtoC11semantics. 5.1.2.4p11 and p16). In the absence of a write instruction, no fence or
AlthoughdesignedforARMv7/POWER,mostofthearguments otheroperationcanpropagateaccumulatedinformationtoanotherthread—
developedintheproofinformallytranslatetotherulesofC11ina inotherwords,itisnotpossibletoestablishahappens-beforepathbetween
tworeadsindifferentthreadswithoutaninterveningwrite.Hencethere-
lianceonseq_cstprimitives,enforcingasequentiallyconsistenttotalorder.
2Henceathreadeventuallyreachesastationarystatewhereb=t;itshould
benotedthatthemodeldoesnotguaranteeprogress;itislegalforathread 4See[6]7.17.3p9.
toenduploopingonanon-finalstatewhereb=tbutb(cid:54)=βnF. 5See[6]7.17.3p11.
6.1 Syntheticbenchmarks In all diagrams, we have included a set of points labeled
nofences,forcomparisonpurposes.Thesecorrespondtotheleast
We designed a synthetic benchmark to simulate the depth-first
commondenominatoramongallthetestedbarrierplacementstrate-
traversalofabalancedtree—withbreadthbanddepthd—ofempty
gies:onlyrelaxedCASoperationsareincluded,withotherwiseno
tasksbyamainworkerthread,reproducingtheprototypicalexecu-
memory barriers. The nofences version violates the semantics of
tionofaCilkprogram.Oneormorethievesattempttostealthese
thework-stealingdeque.Eachofourproposedimplementationsof
tasks.Forrobustnessandpredictability,theworkeralwayscreates
the algorithm can be seen as adding a different set of barriers to
andpushesthesamenumberoftasksinthedeque,followingthe
nofences,makingitcorrect.Hence,resultsobtainedwithnofences
depth-firstpattern,regardlessofwhetheraspecificcontinuationhas
shouldbetakenasnomorethanageneralbaseline,asthecomplete
beenstolenbyanotherthread(itissimplyrecordedasstolen,but
lackoffencescanleadtounexpectedbehavior.Forinstance,Fig-
subsequenttasksspawnnormallyandlocally).Thethievesperform
ures3and4showgreaterthroughputvaluesathighcontentionfor
stealactionsataconfigurablerate,anddiscardstolentasks.
thecomb-shapedworkloadonARM.Thosearetheresultsofalong
Wehaveexperimentedwithtwodifferentmethodsofstealdis-
tailoffastemptytakeoperations,anartifactduetothenatureof
tribution, the goal being to uniformly spread the contention over
thecomb-shapedtestandtheabsenceofsynchronizationbetween
theentirelifeoftheworkerthread.Thefirstmethodisbasedonthe
emptytakeandsteal(enabledbythelackofbarriersintake).8
CPUclockofthecorededicatedtoeachthief;withthistechnique,
theclockisregularlysampledandtheappropriatenumberofsteal b=1;d=107 b=3;d=15
operationsisperformedaccordingly.Thesecondmethodreliesona Corei7(2threads) 4.87862×108 3.60838×108
randomnumbergenerator,calledinabusyloop,whichallowssteal Opteron(2threads) 2.55142×108 2.04978×108
operationswithasetprobability.Whiletheclock-basedapproach Tegra3(2threads) 5.47223×107 4.12112×107
producesmorereliableresults,itcanonlybeusedifalow-overhead Corei7(4threads) 4.88018×108 3.66404×108
CPUclockisavailablefromuserspace,whichisunfortunatelynot Opteron(24threads) 2.56214×108 2.03235×108
thecaseonourARM-basedsystem.6Conversely,thesecondtech- Tegra3(4threads) 5.48473×107 4.11242×107
nique suffers from imprecision when targeting smaller ranges of
frequencies, which is necessary on faster processors or when the Table2. Near-idealthroughput(s−1)
numberofcoresincreases.7 Hence,theformerisusedonx86and
thelatteronARMv7,withappropriateempiricaltuningtogather Alltheresultsbasedonthemixedpushandtake“tree”work-
resultsoveracommonrepresentativerangeofstealthroughput. load show a marked improvement of the hand-written native and
Weselectedtwoworkloads:areasonablybroadtree(b=3;d= c11 versions over the naive sequentially consistent translation of
15) and, as a special case, a degenerate comb-shaped tree (b = the original Chase–Lev algorithm, seqcst. While the relative gain
1;d = 107).Theformerismeanttoreproducenormalcontention remainsstableatalllevelsofcontentionontheCorei7andTegra3,
withstealoperationsalongsidebothpushandtake,whilethelatter itdropssharplyontheOpteron,presumablybecauseofthehigher
illustratesacaseofcontentionbetweentakeandstealonly. numberofcores.Nevertheless,forlowvalues,whichmoreclosely
We measure the time taken by the worker thread to complete modelrealisticscenarios,theoptimizedimplementationsperform
thespecifiednumberoftaskcreationsandconsumptions.Thisin atleast1.5timesbetterthanseqcstonbothx86andARM.
turn serves to compute the push/take throughput—the combined Comparing x86 and ARM, we note that a higher relative
numberofofpushandtakeoperationscompletedperunitoftime, throughputisachievedonARM(peakatabove85%)thanonx86
as well as the effective steal throughput, defined as follows: the (peak at above 50%), indicating that the first serializing instruc-
test protocol strives to perform a number of steals over time, at tionintroducedinthecodeisverycostlyonx86,especiallyifitis
a fixed, nominal steal throughput; the effective throughput is the addedtothecriticalpath(asisthecaseinnative,c11andseqcst,
real throughput as could be observed after the experiment, i.e., butnotinnofences).Thiscouldsuggesteitherthestrongerguaran-
howmanystealswereactuallyperformedduringthelifetimeofthe tees of the x86 memory model—a full memory fence is required
workerthread.Thesemetricsprovideameasureoftheefficiencyof tolinearizehistoryinordertomaintaintotalstoreorder[13]—or
thealgorithmonitscriticalpathatvariouslevelsofcontention. aggressivelocaloptimizationsforsingle-threadexecutionwithout
In order to assess the impact of the added barriers on the dif- communication.
ferentarchitectures,rawthroughputvalueshavebeennormalized Fromtheseobservations,wecanpostulatethatadvancedARM
bythenear-idealthroughputonthesameworkload(seeTable2), architectures such as the Tegra 3 benefit the most from a well-
obtainedonasinglethreadwithnocontentionandnosynchroniza- written concurrent program that takes full advantage of the flexi-
tion: memory barriers are replaced with simple compiler fences, bilityallowedbytheirmemorymodel,andconverselystrugglerel-
andCASoperationswithasimplebranchandconditionalassign- ativelymorewithliteralinterpretationsofalgorithmsdesignedwith
ment.Thesenumbersprovideagoodapproximationoftheupper stricter,simplerhypothesesinmind.
boundontheachievablethroughputoneachmachine,thoughother
6.2 Task-parallelbenchmarks
minorfactorscancontributetohigherobservablevalues.Inpartic-
ular,itshouldbenotedthatcountingthethroughputinnumberof We further experiment on common task-parallel benchmarks,
operationspersecondis,bydesign,ageneralization:theexecution mostly extracted from the Cilk benchmark suite,9 to evaluate the
timeofeachoperationdependsonitsnatureandtheexactcontrol impactofthememorybarrieroptimizationonrealisticworkloads
pathtaken;forexample,aninvocationoftakereturningemptywill andload-balancingscenarios.
befasterthanonereturningatask. Fibonacci isthetree-recursivecomputationofthe35th Fibonacci
number;itillustratestherawcostoftheschedulingalgorithm
aseachtaskonlyperformsasingleaddition.
6TheARMv7C15cyclecounterregistercanonlybequeriediffirstenabled
from kernel mode, and is delegated to a monitoring co-processor, with 8Inthecasewherethedequeisempty,neithertakenorstealneedstoexe-
unclearconsequencesforthebus,caches,andmemorymodelasawhole. cuteaCASinstruction;furthermore,intheabsenceofbarriers,theARMv7
7On higher end processors with multiple cores acting as thieves, higher memorymodeldoesnotrequiresuccessivedecrementsandincrementsof
stealprobabilitiescanyieldmanytimesmorestealattemptsthanthereare bottomintaketopropagatetothethieves.
taskscreatedoverasetperiod. 9http://supertech.csail.mit.edu/cilk