a decision procedure for detecting atomicity violations ...a decision procedure for detecting...

A Decision Procedurefor Detecting Atomicity Violations

for Communicating Processes with Locks?

Nicholas Kidd1, Peter Lammich2, Tayssir Touili3, and Thomas Reps1,4

1 University of Wisconsin, {kidd,reps}@cs.wisc.edu2 Westfalische Wilhelms-Universitat Munster, [email protected]

3 LIAFA, CNRS & Universite Paris Diderot, [email protected] GrammaTech, Inc.

Abstract. We present a new decision procedure for detecting property vi-olations in pushdown models for concurrent programs that use lock-basedsynchronization, where each thread’s lock operations are properly nested(a la synchronized methods in Java). The technique detects violationsexpressed as indexed phase automata (PAs)—a class of non-deterministicfinite automata in which the only loops are self-loops.

Our interest in PAs stems from their ability to capture atomic-set serial-izability violations. (Atomic-set serializability is a relaxation of atomicityto only a user-specified set of memory locations.) We implemented thedecision procedure and applied it to detecting atomic-set-serializabilityviolations in models of concurrent Java programs. Compared with a priormethod based on a semi-decision procedure, not only was the decisionprocedure 7.5X faster overall, but the semi-decision procedure timed outon about 68% of the queries versus 4% for the decision procedure.

1 Introduction

Pushdown systems (PDSs) are a formalism for modeling the interproceduralcontrol flow of recursive programs. Likewise, multi-PDSs have been used to modelthe set of all interleaved executions of a concurrent program with a finite numberof threads [1–7]. This paper presents a decision procedure for multi-PDS modelchecking with respect to properties expressed as indexed phase automata (PAs)—aclass of non-deterministic finite automata in which the only loops are self-loops.The decision procedure handles (i) reentrant locks, (ii) an unbounded number ofcontext switches, and (iii) an unbounded number of lock acquisitions and releasesby each PDS. The decision procedure is compositional : each PDS is analyzedindependently with respect to the PA, and then a single compatibility check isperformed that ties together the results obtained from the different PDSs.

Our interest in PAs stems from their ability to capture atomic-set serializability(AS-serializability) violations. AS-serializability was proposed by Vaziri et al. [8]? Supported by NSF under grants CCF-0540955, CCF-0524051, and CCF-0810053, by

AFRL under contract FA8750-06-C-0249, and by ONR under grant N00014-09-1-0510.

2 Nicholas Kidd, Peter Lammich, Tayssir Touili, and Thomas Reps

as a relaxation of the atomicity property [9] to only a user specified set of fieldsof an object. (A detailed example is given in §2.) In previous work by someof the authors [10], we developed techniques for verifying AS-serializability forconcurrent Java programs. Our tool first abstracts a concurrent Java program intoEML, a modeling language based on multi-PDSs and a finite number of reentrantlocks. The drawback of the approach that we have used to date is that an EMLprogram is compiled into a communicating pushdown system (CPDS) [4, 5], forwhich the required model-checking problem is undecidable. (A semi-decisionprocedure is used in [10].)

Kahlon and Gupta [7] explored the boundary between decidability and unde-cidability for model checking multi-PDSs that synchronize via nested locks. Oneof their results is an algorithm to decide if a multi-PDS satisfies an (indexed)LTL formula that makes use of only atomic propositions, the “next” operatorX, and the “eventually” operator F. In the case of a 2-PDS, the algorithm usesan automaton-pair M = (A,B) to represent a set of configurations of a 2-PDS,where an automaton encodes the configurations of a single PDS in the usualway [11, 12]. For a given logical formula, the Kahlon-Gupta algorithm is definedinductively: from an automaton-pair that satisfies a subformula, they define analgorithm that computes a new automaton-pair for a larger formula that has oneadditional (outermost) temporal operator.

We observed that PAs can be compiled into an LTL formula that uses onlythe X and F operators. (An algorithm to perform the encoding is given in [13,App. A].) Furthermore, [14] presents a sound and precise technique that uses onlynon-reentrant locks to model EML’s reentrant locks. Thus, combining previouswork [10, 14] with the Kahlon-Gupta algorithm provides a decision procedure forverifying AS-serializability of concurrent Java programs!

(Briefly, the technique for replacing reentrant locks with non-reentrant lockspushes a special marker onto the stack the first time a lock is acquired, andrecords the acquisition in a PDS’s state space. All subsequent lock acquires andtheir matching releases do not change the state of the lock or the PDS. Onlywhen the special marker is seen again is the lock then released. This techniquerequires that lock acquisition and releases be properly scoped, which is satisfiedby Java’s synchronized blocks. Consequently, we consider only non-reentrantlocks in the remainder of the paper.)

Unfortunately, [7] erroneously claims that the disjunction operation distributesacross two automaton-pairs. That is, for automaton-pairs M1 = (A1, B1) andM2 = (A2, B2), they claim that the following holds: M1∨M2 = (A1∨A2, B1∨B2).This is invalid because cross-terms arise when attempting to distribute thedisjunct. For example, if B1 ∩ B2 = ∅, then there can be configurations of theform 〈a1 ∈ A1, b2 ∈ B2〉 that would be accepted by (A1∨A2, B1∨B2) but shouldnot be in M1 ∨M2.

To handle this issue properly, a corrected algorithm must use a set ofautomaton-pairs instead of single automaton-pair to represent a set of con-figurations of a 2-PDS.5 Because the size of the set is exponential in the number

5 We confirmed these observations in e-mail with Kahlon and Gupta [15].

A Decision Procedure 3

of locks, in the worst case, their algorithm may perform an exponential numberof individual reachability queries to handle one temporal operator. Furthermore,once reachability from one automaton-pair has been performed, the resultingautomaton pair must be split into a set so that incompatible configurations areeliminated. Thus, it is not immediately clear if the (corrected) Kahlon-Guptaalgorithm is amenable to an implementation that would be usable in practice.

This paper presents a new decision procedure for checking properties specifiedas PAs on multi-PDSs that synchronize via nested locks.6 Unlike the (corrected)Kahlon-Gupta algorithm, our decision procedure uses only one reachability queryper PDS. The key is to use tuples of lock histories (§5): moving from the lockhistories used by Kahlon and Gupta to tuples-of-lock histories introduces amechanism to maintain the correlations between the intermediate configurations.Hence, our decision procedure is able to make use of only a single compatibilitycheck over the tuples-of-lock histories that our analysis obtains for each PDS.The benefit of this approach is shown in the following table, where Procs denotesthe number of processes, L denotes the number of locks, and |PA| denotes thenumber of states in property automaton PA:

PDS State Space Queries

Kahlon-Gupta [7] (corrected) O(2L) O(|PA| · Procs · 2L)

This paper (§6) O(|PA| · 2L) Procs

Because our algorithm isolates the exponential cost in the PDS state space, thatcost can often be side-stepped using symbolic techniques, such as BDDs, asexplained in §7.

This paper makes the following contributions:

– We define a decision procedure for multi-PDS model-checking for PAs. Thedecision procedure handles (i) reentrant locks, (ii) an unbounded number ofcontext switches, (iii) an unbounded number of lock acquisitions and releasesby each PDS, and (iv) a bounded number of phase transitions.

– The decision procedure is compositional : each PDS is analyzed independentlywith respect to the PA, and then a single compatibility check is performedthat ties together the results obtained from the different PDSs.

– We leverage the special form of PAs to give a symbolic implementation thatis more space-efficient than standard BDD-based techniques for PDSs [16].

– We used the decision procedure to detect AS-serializability violations inautomatically-generated models of four programs from the ConTest bench-mark suite [17], and obtained substantially better performance than a priormethod based on a semi-decision procedure [10].

The rest of the paper is organized as follows: §2 provides motivation. §3 definesmulti-PDSs and PAs. §4 reviews Kahlon and Gupta’s decomposition result. §5presents lock histories. §6 presents the decision procedure. §7 describes a symbolicimplementation. §8 presents experimental results. §9 describes related work.

6 We do not consider multi-PDSs that use wait-notify synchronization because reacha-bility analysis of multi-PDSs with wait-notify is undecidable [7].


class Stack {Object[] storage = new Object[10];int item = -1;public static Stack makeStack(){

return new Stack();}public synchronized Object pop(){

Object res = storage[item];storage[item--] = null;return res;}public synchronized void push(Object o){

storage[++item] = o;}public synchronized boolean empty(){

return (item == -1);}}

class Client {// @atomicpublic synchronized Object get(Stack s){

if(!s.empty()) { return s.pop(); }else return null;}public static Client makeClient(){

return new Client();}public static void main(String[] args){

Stack stack = Stack.makeStack();stack.push(new Integer(1));Client client1 = makeClient();Client client2 = makeClient();new Thread("1") { client1.get(stack); }new Thread("2") { client2.get(stack); }}}

1:

get()z }| {Abeg1

empty()z }| {(sR1(i))s . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . .

pop()z }| {(sR1(i)R1(s)R1(i)W1(s)W1(i))s Aend1

2: . . . . . . . . . . . . . .

get()z }| {Abeg2 (sR2(i))s| {z }

empty()

(sR2(i)R2(s)R2(i)W2(s)W2(i))s| {z }pop()

Aend2 . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . .

Fig. 1. Example program and problematic interleaving that violates atomic-set serializ-ability. R and W denote a read and write access, respectively. i and s denote fields itemand storage, respectively. Abeg and Aend denote the beginning and end, respectively,of an atomic code block. The subscripts “1” and “2” are thread ids. “(s” and “)s” denotethe acquire and release operations, respectively, of the lock of Stack stack.

2 Motivating Example

Fig. 1 shows a simple Java implementation of a stack. Class Client is a testharness that performs concurrent accesses on a single stack. Client.get() usesthe keyword “synchronized” to protect against concurrent calls on the sameClient object. The annotation “@atomic” on Client.get() specifies that theprogrammer intends for Client.get() to be executed atomically.

The program’s synchronization actions do not ensure this, however. The rootcause is that the wrong object is used for synchronization: parameter “Stacks” of Client.get() should have been used, instead of Client.get()’s implicitthis parameter. This mistake permits the interleaved execution shown at thebottom of Fig. 1, which would result in an exception being thrown.

This is an example of an atomic-set serializability (AS-serializability)—arelaxation of atomicity [9] to only a specified set of shared-memory locations—violation [8] with respect to s.item and s.storage. AS-serializability violationscan be completely characterized by a set of fourteen problematic access patterns


[8].7 Each problematic pattern is a finite sequence of reads and writes by twothreads to one or two shared memory locations. For the program in Fig. 1 andproblematic pattern “Abeg1;R1(i);W2(s);W2(i);R1(s)”, the accesses that matchthe pattern are underlined in the interleaving shown at the bottom of Fig. 1.

W2(s) W2(i) R1(s)R1(i)Abeg1q1 q2 q3 q4 q5 q6

Σ ΛΛΛΛ

Fig. 2. The PA that accepts the problematic

access pattern in the program from Fig. 1. Σ

is the set of all actions, and Λ is Σ \ {Aend1}.

The fourteen problematic accesspatterns can be encoded as an in-dexed phase automaton (PA). ThePA that captures the problematic ac-cesses of Fig. 1 is shown in Fig. 2. Itsstates—which represent the phasesthat the automaton passes throughto accept a string—are chained to-gether by phase transitions; eachstate has a self-loop for symbols thatcause the automaton to not changestate. (“Indexed” refers to the fact that the index of the thread performing anaction is included in the label of each transition.)

The PA in Fig. 2 “guesses” when a violation occurs. That is, when it observesthat thread 1 enters an atomic code block, such as get(), the atomic-code-block-begin action Abeg1 causes it either to transition to state q2 (i.e., to start the nextphase), or to follow the self-loop and remain in q1. This process continues until itreaches the accepting state. Note that the only transition that allows thread 1to exit an atomic code block (Aend1) is the self-loop on the initial state. Thus,incorrect guesses cause the PA in Fig. 2 to become “stuck” in one of the statesq1 . . . q5 and not reach final state q6.

3 Program Model and Property Specifications

Definition 1. A (labeled) pushdown system (PDS) is a tuple P =(P,Act, Γ,∆, c0), where P is a finite set of control states, Act is a finite setof actions, Γ is a finite stack alphabet, and ∆ ⊆ (P × Γ )×Act× (P × Γ ∗) is afinite set of rules. A rule r ∈ ∆ is denoted by 〈p, γ〉 a

↪−→〈p′, u′〉. A PDS configu-ration 〈p, u〉 is a control state along with a stack, where p ∈ P and u ∈ Γ ∗, andc0 = 〈p0, γ0〉 is the initial configuration. ∆ defines a transition system over the setof all configurations. From c = 〈p, γu〉, P can make a transition to c′ = 〈p′, u′u〉on action a, denoted by c a−→ c′, if there exists a rule 〈p, γ〉 a

↪−→〈p′, u′〉 ∈ ∆. Forw ∈ Act∗, c w−−→ c′ is defined in the usual way. For a rule r = 〈p, γ〉 a

↪−→〈p′, u′〉,act(r) denotes r’s action label a.

A multi-PDS consists of a finite number of PDSs P1, . . . ,Pn that synchronizevia a finite set of locks Locks = {l1, . . . , lL} (i.e., L = |Locks|). The actions Act

7 This result relies on an assumption that programs do not always satisfy: an atomiccode section that writes to one member of a set of correlated locations writes to alllocations in that set (e.g., item and storage of Stack s).


of each PDS consist of lock-acquires (“(i”) and releases (“)i”) for 1 ≤ i ≤ L, plussymbols from Σ, a finite alphabet of non-parenthesis symbols.

The intention is that each PDS models a thread, and lock-acquire and releaseactions serve as synchronization primitives that constrain the behavior of themulti-PDS. We assume that locks are acquired and released in a well-nestedfashion; i.e., locks are released in the opposite order in which they are acquired.

The choice of what actions appear in Σ depends on the intended application.For verifying AS-serializability (see §2 and §7), Σ consists of actions to read andwrite a shared-memory location m (denoted by R(m) and W (m), respectively),and to enter and exit an atomic code section (Abeg and Aend, respectively).

Formally, a program model is a tuple Π = (P1, . . . ,Pn, Locks, Σ). A globalconfiguration g = (c1, . . . , cn, o1, . . . , oL) is a tuple consisting of a local configura-tion ci for each PDS Pi and a valuation that indicates the owner of each lock:for each 1 ≤ i ≤ L, oi ∈ {⊥, 1, . . . , n} indicates the identity of the PDS thatholds lock li. The value ⊥ signifies that a lock is currently not held by any PDS.The initial global configuration is g0 = (c10, . . . , c

n0 ,⊥, . . . ,⊥). A global config-

uration g = (c1, c2, . . . , cn, o1, . . . , oL) can make a transition to another globalconfiguration g′ = (c′1, c2, . . . , cn, o

′1, . . . , o

′L) under the following conditions:

– If c1a−→ c′1 and a /∈ {(i, )i}, then g′ = (c′1, c2, . . . , cn, o1, . . . , oL).

– If c1(i−−→ c′1 and g = (c1, c2, . . . , cn, o1, . . . , oi−1,⊥, oi+1, . . . , oL), then g′ =

(c′1, c2, . . . , cn, o1, . . . , oi−1, 1, oi+1, . . . , oL).– If c1

)i−−→ c′1 and g = (c1, c2, . . . , cn, o1, . . . , oi−1, 1, oi+1, . . . , oL), then g′ =(c′1, c2, . . . , cn, o1, . . . , oi−1,⊥, oi+1, . . . , oL).

For 1 < j ≤ n, a global configuration (c1, . . . , cj , . . . , cn, o1, . . . , oL) can make atransition to (c1, . . . , c′j , . . . , cn, o

′1, . . . , o

′L) in a similar fashion.

A program property is specified as an indexed phase automaton.

Definition 2. An indexed phase automaton (PA) is a tuple (Q, Id, Σ, δ),where Q is a finite, totally ordered set of states {q1, . . . , q|Q|}, Id is a finite set ofthread identifiers, Σ is a finite alphabet, and δ ⊆ Q× Id×Σ ×Q is a transitionrelation. The transition relation δ is restricted to respect the order on states: foreach transition (qx, i, a, qy) ∈ δ, either y = x or y = x+ 1. We call a transitionof the form (qx, i, a, qx+1) a phase transition. The initial state is q1, and thefinal state is q|Q|.

The restriction on δ in Defn. 2 ensures that the only loops in a PA are “self-loops” on states. We assume that for every x, 1 ≤ x < |Q|, there is only onephase transition of the form (qx, i, a, qx+1) ∈ δ. (A PA that has multiple suchtransitions can be factored into a set of PAs, each of which satisfy this property.)Finally, we only consider PAs that recognize a non-empty language, which meansthat a PA must have exactly (|Q| − 1) phase transitions.

For the rest of this paper we consider 2-PDSs, and fix Π = (P1,P2, Locks, Σ)and A = (Q, Id, Σ, δ); however, the techniques easily generalize to multi-PDSs(see [13, App. B]), and our implementation is for the generic case. Given Π andA, the model-checking problem of interest is to determine if there is an executionthat begins at the initial global configuration g0 that drives A to its final state.


4 Path Incompatibility

The decision procedure analyzes the PDSs of Π independently, and then checksif there exists a run from each PDS that can be performed in interleaved parallelfashion under the lock-constrained transitions of Π. To do this, it makes use ofa decomposition result, due to Kahlon and Gupta [7, Thm. 1], which we nowreview.

Suppose that PDS Pk, for k ∈ {1, 2}, when started in (single-PDS) config-uration ck and executed alone, is able to reach configuration c′k using the rulesequence ρk. Let LocksHeld(Pk, (b1, b2, o1, . . . , oL)) denote {li | oi = k}; i.e., theset of locks held by PDS Pk at global configuration (b1, b2, o1, . . . , oL).

Along a rule sequence ρk and for an initially-held lock li and finally-held locklf , we say that the initial release of li is the first release of li, and that the finalacquisition of lf is the last acquisition of lf . Note that for execution to proceedalong ρk, Pk must hold an initial set of locks at ck that is a superset of the set ofinitial releases along ρk; i.e., not all initially-held locks need be released. Similarly,Pk’s final set of locks at c′k must be a superset of the set of final acquisitionsalong ρk.

Theorem 1. (Decomposition Theorem [7].) Suppose that PDS Pk, whenstarted in configuration ck and executed alone, is able to reach configurationc′k using the rule sequence ρk. For Π = (P1,P2, Locks, Σ), there does not existan interleaving of paths ρ1 and ρ2 from global configuration (c1, c2, o1, . . . , oL)to global configuration (c′1, c

′2, o′1, . . . , o

′L) iff one or more of the following five

conditions hold:

1. LocksHeld(P1, (c1, c2, o1, . . . , oL)) ∩ LocksHeld(P2, (c1, c2, o1, . . . , oL)) 6= ∅2. LocksHeld(P1, (c′1, c

′2, o′1, . . . , o

′L)) ∩ LocksHeld(P2, (c′1, c

′2, o′1, . . . , o

′L)) 6= ∅

3. In ρ1, P1 releases lock li before it initially releases lock lj, and in ρ2, P2

releases lj before it initially releases lock li.4. In ρ1, P1 acquires lock li after its final acquisition of lock lj, and in ρ2, P2

acquires lock lj after its final acquisition of lock li,5. (a) In ρ1, P1 acquires or uses a lock that is held by P2 throughout ρ2, or

(b) in ρ2, P2 acquires or uses a lock that is held by P1 throughout ρ1.

Intuitively, items 3 and 4 capture cycles in the dependence graph of lockoperations: a cycle is a proof that there does not exist any interleaving of rulesequences ρ1 and ρ2 that adheres to the lock-constrained semantics of Π. Ifthere is a cycle, then ρ1 (ρ2) can complete execution but not ρ2 (ρ1), or neithercan complete because of a deadlock. The remaining items model standard locksemantics: only one thread may hold a lock at a given time.

5 Extracting Information from PDS Rule Sequences

To employ Thm. 1, we now develop methods to extract relevant informationfrom a rule sequence ρk for PDS Pk. As in many program-analysis problems that


involve matched operations [18]—in our case, lock-acquire and lock-release—it isuseful to consider semi-Dyck languages [19]: languages of matched parenthesesin which each parenthesis symbol is one-sided . That is, the symbols “(” and “)”match in the string “()”, but do not match in “)(”.8

LetΣ be a finite alphabet of non-parenthesis symbols. The semi-Dyck languageof well-balanced parentheses over Σ ∪ {(i, )i | 1 ≤ i ≤ L} can be defined by thefollowing context-free grammar, where e denotes a member of Σ:

matched → ε | e matched | (i matched )i matched [for 1 ≤ i ≤ L]Because we are interested in paths that can begin and end while holding a set oflocks, we define the following partially-matched parenthesis languages:

unbalR → ε | unbalR matched )i unbalL → ε | (i matched unbalLThe language of words that are possibly unbalanced on each end is defined by

suffixPrefix → unbalR matched unbalL

Example 1. Consider the following suffixPrefix string, in which the positionsbetween symbols are marked A–W. Its unbalR, matched, and unbalL componentsare the substrings A–N, N–P, and P–W, respectively.

A

)1B

(2C

)2D

)3E

(2F

(4G

(5H

)5I

)4J

(6K

)6L

)2M

)7N

(6O

)6P

(4Q

(2R

)2S

(2T

(7U

)7V

(8W

Let wk ∈ L(suffixPrefix) be the word formed by concatenating the actionsymbols of the rule sequence ρk. One can see that to use Thm. 1, we merelyneed to extract the relevant information from wk. That is, items 3 and 4 requireextracting (or recording) information from the unbalR and unbalL portions of wk,respectively; item 5 requires extracting information from the matched portion ofwk; and items 1 and 2 require extracting information from the initial and finalparse configurations of wk.

The information is obtained using acquisition histories (AH) and releasehistories (RH) for locks, as well as ρk’s release set (R), use set (U), acquisitionset (A), and held-throughout set (HT).

– The acquisition history (AH) [7] for a finally-held lock li is the union of theset {li} with the set of locks that are acquired (or acquired and released)after the final acquisition of li.9

– The release history (RH) [7] of an initially-held lock li is the union of the set{li} with the set of locks that are released (or acquired and released) beforethe initial release of li.

– The release set (R) is the set of initially-released locks.– The use set (U) is the set of locks that form the matched part of wk.– The acquisition set (A) is the set of finally-acquired locks.– The held-throughout set (HT) is the set of initially-held locks that are not

released.8 The language of interest is in fact regular because the locks are non-reentrant.

However, the semi-Dyck formulation provides insight into how one extracts therelevant information from a rule sequence.

9 This is a slight variation from [7]; we include li in the acquisition history of lock li.


A lock history is a six-tuple (R, RH,U, AH,A,HT), where R, U, A, and HT

are the sets defined above, and AH (RH) is a tuple of L acquisition (release)histories, one for each lock li, 1 ≤ i ≤ L. Let ρ = [r1, . . . , rn] be a rule sequencethat drives a PDS from some starting configuration to an ending configuration,and let I be the set of locks held at the beginning of ρ. We define an abstractionfunction η(ρ, I) from rule sequences and initially-held locks to lock histories;η(ρ, I) uses an auxiliary function, post, which tracks R, RH, U, AH, A, and HTfor each successively longer prefix.

η([], I) = (∅, ∅L, ∅, ∅L, ∅, I)η([r1, . . . , rn], I) = post(η([r1, . . . , rn−1], I), act(rn)), where

post((R, cRH,U, cAH,A,HT), a) =8>>>>>>>>>>>>>>>>>><>>>>>>>>>>>>>>>>>>:

(R, cRH,U, cAH,A,HT) if a /∈ {(i, )i}(R, cRH,U, cAH

′,A ∪ {li},HT) if a = (i

where cAH′[j] =

8<:{li} if j = i∅ if j 6= i and lj /∈ AcAH[j] ∪ {li} if j 6= i and lj ∈ A

(R, cRH,U ∪ {li}, cAH′,A\{li},HT\{li}) if a = )i and li ∈ A

where cAH′[j] =

∅ if j = icAH[j] otherwise

(R ∪ {li}, cRH′,U, cAH,A,HT\{li}) if a = )i and li /∈ A

where cRH′[j] =

{li} ∪ U ∪ R if j = icRH[j] otherwise

Example 2. Suppose that ρ is a rule sequence whose labels spell out the stringfrom Example 1, and I = {1, 3, 7, 9}. Then η(ρ, I) returns the following lockhistory (only lock indices are written):

({1, 3, 7}, 〈{1}, ∅, {1, 2, 3}, ∅, ∅, ∅, {1, 2, 3, 4, 5, 6, 7}, ∅, ∅〉,{6}, 〈∅, {2, 7, 8}, ∅, {2, 4, 7, 8}, ∅, ∅, ∅, {8}, ∅ 〉, {2, 4, 8}, {9}).

Note: R and A are included above only for clarity; they can be recovered fromRH and AH, as follows: R = {i | RH[i] 6= ∅} and A = {i | AH[i] 6= ∅}. In addition,from LH = (R, RH,U, AH,A,HT), it is easy to see that the set I of initially-heldlocks is equal to (R ∪ HT), and the set of finally-held locks is equal to (A ∪ HT).

Definition 3. Lock histories LH1 = (R1, RH1, U1, AH1, A1, HT1) and LH2 =(R2, RH2, U2, AH2, A2, HT2) are compatible, denoted by Compatible(LH1, LH2),iff all of the following five conditions hold:

1.(R1 ∪ HT1) ∩ (R2 ∪ HT2) = ∅ 2.(A1 ∪ HT1) ∩ (A2 ∪ HT2) = ∅3. 6 ∃ i, j . lj ∈ cAH1[i] ∧ li ∈ cAH2[j] 4. 6 ∃ i, j . lj ∈ cRH1[i] ∧ li ∈ cRH2[j]

5.(A1 ∪ U1) ∩ HT2 = ∅ ∧ (A2 ∪ U2) ∩ HT1 = ∅

Each conjunct verifies the absence of the corresponding incompatibility conditionfrom Thm. 1: conditions 1 and 2 verify that the initially-held and finally-heldlocks of ρ1 and ρ2 are disjoint, respectively; conditions 3 and 4 verify the absenceof cycles in the acquisition and release histories, respectively; and condition 5verifies that ρ1 does not use a lock that is held throughout in ρ2, and vice versa.


Abeg1 R1(i) W2(s) W2(i) R1(s)1:q3 q4 q4 q5q1 q2 q2 q3 q5 q6

Abeg1 R1(i) W2(s) W2(i) R1(s)2:q1 q2 q2 q3 q5 q6q3 q4 q4 q5

Abeg1 R1(i) W2(s) W2(i) R1(s)q3 q4 q4 q5q1 q2 q2 q3 q5 q6Π:

Fig. 3. Π: bad interleaving of Fig. 2, showing only the actions that cause a phasetransition. 1: the same interleaving from Thread 1’s point of view. The dashed boxesshow where Thread 1 guesses that Thread 2 causes a phase transition. 2: the same butfrom Thread 2’s point of view and with the appropriate guesses.

6 The Decision Procedure

As noted in §4, the decision procedure analyzes the PDSs independently. Thisdecoupling of the PDSs has two consequences.

First, when P1 and A are considered together, independently of P2, theycannot directly “observe” the actions of P2 that cause A to take certain phasetransitions. Thus, P1 must guess when P2 causes a phase transition, and viceversa for P2. An example of the guessing is shown in Fig. 3. The interleavinglabeled “Π” is the bad interleaving from Fig. 2, but focuses on only the PDSactions that cause phase transitions. The interleaving labeled “1” shows, via thedashed boxes, where P1 guesses that P2 caused a phase transition. Similarly, theinterleaving labeled “2” shows the guesses that P2 must make.

Second, a post-processing step must be performed to ensure that only thosebehaviors that are consistent with the lock-constrained behaviors of Π areconsidered. For example, if P1 guesses that P2 performs the W2(s) action tomake the PA transition from state q3 to state q4 (the dashed box for interleaving “1”in Fig. 3) while it is still executing the empty() method (see Fig. 2), the behavioris inconsistent with the semantics of Π. This is because both threads would holdthe lock associated with the shared “Stack s” object. The post-processing stepensures that such behaviors are not allowed.

6.1 Combining a PDS with a PA

To define a compositional algorithm, we must be able to analyze P1 and Aindependently of P2, and likewise for P2 and A. Our approach is to combine Aand P1 to define a new PDS PA1 using a cross-product-like construction. Themain difference is that lock histories and lock-history updates are incorporatedin the construction.

Recall that the goal is to determine if there exists an execution of Π that drivesA to its final state. Any such execution must make |Q|−1 phase transitions. Hence,


a valid interleaved execution must be able to reach |Q| global configurations, onefor each of the |Q| phases.

Lock histories encode the constraints that a PDS path places on the set ofpossible interleaved executions of Π. A desired path of an individual PDS mustalso make |Q| − 1 phase transitions, and hence our algorithm keeps track of |Q|lock histories, one for each phase. This is accomplished by encoding into the statespace of PA1 a tuple of |Q| lock histories. A tuple maintains the sequence of lockhistories for one or more paths taken through a sequence of phases. In addition,a tuple maintains the correlation between the lock histories of each phase, whichis necessary to ensure that only valid executions are considered. The rules of PA1are then defined to update the lock-history tuple accordingly. The lock-historytuples are used later to check whether some scheduling of an execution of Π canactually perform all of the required phase transitions.

Let LH denote the set of all lock histories, and let LH = LH|Q| denote theset of all tuples of lock histories of length |Q|. We denote a typical lock history byLH, and a typical tuple of lock histories by LH. LH[i] denotes the ith componentof LH.

Our construction makes use of the phase-transition function on LHs definedas follows: ptrans((R, RH,U, AH,A,HT)) = (∅, ∅L, ∅, ∅L, ∅,A ∪ HT). This functionis used to encode the start of a new phase: the set of initially-held locks is theset of locks held at the end of the previous phase.

Let Pi = (Pi,Acti, Γi, ∆i, 〈p0, γ0〉) be a PDS, Locks be a set of locks of sizeL, A = (Q, Id,Σ, δ) be a PA, and LH be a tuple of lock histories of length |Q|.We define the PDS PAi = (PAi , ∅, Γi, ∆Ai , 〈pA0 , γ0〉), where PAi ⊆ Pi × Q × LH.The initial control state is pA0 = (p0, q1, LH∅), where LH∅ is the lock-history tuple(∅, ∅L, ∅, ∅L, ∅, ∅)|Q|. Each rule r ∈ ∆Ai performs only a single update to the tupleLH, at an index x determined by a transition in δ. The update is denoted byLH[x 7→ e], where e evaluates to an LH. Two kinds of rules are introduced toaccount for whether a transition in δ is a phase transition or not:

1. Non-phase Transitions: LH′

= LH[x 7→ post(LH[x], a)].

(a) For each rule 〈p, γ〉 a↪−→〈p′, u〉 ∈ ∆i and transition (qx, i, a, qx) ∈ δ, there

is a rule r = 〈(p, qx, LH), γ〉 ↪−→〈(p′, qx, LH′), u〉 ∈ ∆Ai .

(b) For each rule 〈p, γ〉 a↪−→〈p′, u〉 ∈ ∆i, a ∈ {(k, )k}, and each qx ∈ Q, there

is a rule r = 〈(p, qx, LH), γ〉 ↪−→〈(p′, qx, LH′), u〉 ∈ ∆Ai .

2. Phase Transitions: LH′

= LH[(x+ 1) 7→ ptrans(LH[x])].

(a) For each rule 〈p, γ〉 a↪−→〈p′, u〉 ∈ ∆i and transition (qx, i, a, qx+1) ∈ δ,

there is a rule r = 〈(p, qx, LH), γ〉 ↪−→〈(p′, qx+1, LH′), u〉 ∈ ∆Ai .

(b) For each transition (qx, j, a, qx+1) ∈ δ, j 6= i, and for each p ∈ Pi andγ ∈ Γi, there is a rule r = 〈(p, qx, LH), γ〉 ↪−→〈(p, qx+1, LH

′), γ〉 ∈ ∆A.


input : A 2-PDS Π = (P1,P2, Locks, Σ) and a PA A.output: true if Π can drive A to its final state.let A1

post∗ ← post∗PA1; let A2

post∗ ← post∗PA2;1

foreach p1 ∈ P1, cLH1 s.t. ∃u1 ∈ Γ ∗1 : 〈(p1, q|Q|,cLH1), u1〉 ∈ L(A1post∗) do2

foreach p2 ∈ P2, cLH2 s.t. ∃u2 ∈ Γ ∗2 : 〈(p2, q|Q|,cLH2), u2〉 ∈ L(A2post∗) do3

if Compatible(cLH1,cLH2) then4

return true;5

return false;6

Algorithm 1: The decision procedure. The two tests of the form “∃uk ∈ Γ ∗k :

〈(pk, q|Q|,cLHk), uk〉 ∈ L(Akpost∗)” can be performed by finding any path in Ak

post∗ from

state (pk, q|Q|,cLHk) to the final state.

Rules defined by item 1(a) make sure that PAi is constrained to follow the self-loopson PA state qx. Rules defined by item 1(b) allow for PAi to perform lock acquiresand releases. Recall that the language of a PA is only over the non-parenthesisalphabet Σ, and does not constrain the locking behavior. Consequently, a phasetransition cannot occur when PAi is acquiring or releasing a lock. Rules definedby item 2(a) handle phase transitions caused by PAi . Finally, rules defined byitem 2(b) implement PAi ’s guessing that another PDS PAj , j 6= i, causes a phasetransition, in which case PAi has to move to the next phase as well.

6.2 Checking Path Compatibility

For a generated PDS PAk , we are interested in the set of paths that begin inthe initial configuration 〈pA0 , γ0〉 and drive A to its final state q|Q|. Each suchpath ends in some configuration 〈(pk, q|Q|, LHk), u〉, where u ∈ Γ ∗. Let ρ1 andρ2 be such paths from PA1 and PA2 , respectively. To determine if there exists acompatible scheduling for ρ1 and ρ2, we use Thm. 1 on each component of thelock-history tuples LH1 and LH2 from the ending configurations of ρ1 and ρ2:

Compatible(LH1, LH2)⇐⇒|Q|∧i=1

Compatible(LH1[i], LH2[i]).

Due to recursion, PA1 and PA2 could each have an infinite number of such paths.However, each path is abstracted as a tuple of lock histories LH, and there areonly a finite number of tuples in LH; thus, we only have to check a finite numberof (LH1, LH2) pairs. For each PDS PA = (PA,Act, Γ,∆, cA0 ), we can identifythe set of relevant LH tuples by computing the set of all configurations thatare reachable starting from the initial configuration, post∗PA(cA0 ), using standardautomata-based PDS techniques [11, 12]. (Because the initial configuration isdefined by the PDS PA, henceforth, we merely write post∗PA .) That is, becausethe construction of PA removed all labels, we can create a P-(multi)-automaton[11] Apost∗ that accepts exactly the set of configurations post∗PA .

Alg. 1 gives the algorithm to check whether Π can drive A to its final state.


Theorem 2. For 2-PDS Π = (P1,P2, Locks, Σ) and PA A, there exists anexecution of Π that drives A to its final state iff Alg. 1 returns true.

Proof. See [13, App. D.1]. ut

7 A Symbolic Implementation

Alg. 1 solves the multi-PDS model-checking problem for PAs. However, an imple-mentation based on symbolic techniques is required because it would be infeasibleto perform the final explicit enumeration step specified in Alg. 1, lines 2–5. Onepossibility is to use Schwoon’s BDD-based PDS techniques [16]; these representthe transitions of a PDS’s control-state from one configuration to another as arelation, using BDDs. This approach would work with relations over Q × LH,which requires using |Q|2|LH|2 BDD variables, where |LH| = 2L+ 2L2.

This section describes a more economical encoding that needs only (|Q|+1)|LH|BDD variables. Our approach leverages the fact that when a property is specifiedwith a phase automaton, once a PDS makes a phase transition from qx to qx+1,the first x entries in LH tuples are no longer subject to change. In this situation,Schwoon’s encoding contains redundant information; our technique eliminatesthis redundancy.

We explain the more economical approach by defining a suitable weightdomain for use with a weighted PDS (WPDS) [4, 20]. A WPDS W = (P,S, f)is a PDS P = (P,Act, Γ,∆, c0) augmented with a bounded idempotent semiringS = (D,⊗,⊕, 1, 0) (see [13, App. C]), and a function f : ∆→ D that assigns asemiring element d ∈ D to each rule r ∈ ∆. When working with WPDSs, theresult of a post∗ computation is a weighted automaton. For the purposes of thispaper, we view the weighted automaton Apost∗ = post∗W as a function from aregular set of configurations C to the sum-over-all-paths from c0 to all c ∈ C;i.e., Apost∗(C) =

⊕{v | ∃c ∈ C : c0

r1...rn−−−−−→ c, v = f(r1) ⊗ . . . ⊗ f(rn)}, wherer1 . . . rn is a sequence of rules that transforms c0 into c. For efficient algorithmsfor computing both Apost∗ and Apost∗(C), see [4, 20].

Definition 4. Let S be a finite set; let A ⊆ Sm+1 and B ⊆ Sp+1 be relations ofarity m+ 1 and p+ 1, respectively. The generalized relational compositionof A and B, denoted by “A ; B”, is the following subset of Sm+p:A ; B = {〈a1, . . . , am, b2, . . . , bp+1〉 | 〈a1, . . . , am, x〉 ∈ A∧〈x, b2, . . . , bp+1〉 ∈ B}.

Definition 5. Let S be a finite set, and θ be the maximum number of phases ofinterest. The set of all θ-term formal power series over z, with relation-valued coefficients of different arities, is

RFPS[S, θ] = {∑θ−1i=0 ciz

i | ci ⊆ Si+2}.A monomial is written as cizi (all other coefficients are understood to be ∅);a monomial c0z0 denotes a constant. The multi-arity relational weightdomain over S and θ is defined by (RFPS[S, θ],×,+, Id, ∅), where × is poly-nomial multiplication in which generalized relational composition and ∪ are


used to multiply and add coefficients, respectively, and terms cjzj for j ≥ θ aredropped; + is polynomial addition using ∪ to add coefficients; Id is the constant{〈s, s〉 | s ∈ S}z0; and ∅ is the constant ∅z0.

We now define the WPDS Wi = (PWi ,S, f) that results from taking the prod-uct of PDS Pi = (Pi,Acti, Γi, ∆i, 〈p0, γ0〉) and phase automatonA = (Q, Id,Σ, δ).The construction is similar to that in §6.1, i.e., a cross product is performed thatpairs the control states of Pi with the state space of A. The difference is that thelock-history tuples are removed from the control state, and instead are modeledby S, the multi-arity relational weight domain over the finite set LH and θ = |Q|.We define PWi = (Pi ×Q, ∅, Γi, ∆Wi , 〈(p0, q1), γ0〉), where ∆Wi and f are definedas follows:

1. Non-phase Transitions: f(r) = {〈LH1, LH2〉 | LH2 = post(LH1, a)}z0.

(a) For each rule 〈p, γ〉 a↪−→〈p′, u〉 ∈ ∆i and transition (qx, i, a, qx) ∈ δ, there

is a rule r = 〈(p, qx), γ〉 ↪−→〈(p′, qx), u〉 ∈ ∆Wi .(b) For each rule 〈p, γ〉 a

↪−→〈p′, u〉 ∈ ∆i, a ∈ {(k, )k}, and for each qx ∈ Q,there is a rule r = 〈(p, qx), γ〉 ↪−→〈(p′, qx), u〉 ∈ ∆Wi .

2. Phase Transitions: f(r) = {〈LH, LH, ptrans(LH)〉 | LH ∈ LH}z1.

(a) For each rule 〈p, γ〉 a↪−→〈p′, u〉 ∈ ∆i and transition (qx, i, a, qx+1) ∈ δ,

there is a rule r = 〈(p, qx), γ〉 ↪−→〈(p′, qx+1), u〉 ∈ ∆Wi .(b) For each transition (qx, j, a, qx+1) ∈ δ, j 6= i, and for each p ∈ Pi and

γ ∈ Γi, there is a rule r = 〈(p, qx), γ〉 ↪−→〈(p, qx+1), γ〉 ∈ ∆W .

A multi-arity relational weight domain is parameterized by the quantityθ—the maximum number of phases of interest—which we have picked to be |Q|.We must argue that weight operations performed during model checking do notcause this threshold to be exceeded. For configuration 〈(p, qx), u〉 to be reachablefrom the initial configuration 〈(p0, q1), γ0〉 of some WPDS Wi, PA A must makea sequence of transitions from states q1 to qx, which means that A goes throughexactly x− 1 phase transitions. Each phase transition multiplies by a weight ofthe form c1z

1; hence, the weight returned by Apost∗({〈(p, qx), u〉}) is a monomialof the form cx−1z

x−1. The maximum number of phases in a PA is |Q|, and thusthe highest-power monomial that arises is of the form c|Q|−1z

|Q|−1. (Moreover,during post∗Wk

as computed by the algorithm from [20], only monomial-valuedweights ever arise.)

Alg. 2 states the algorithm for solving the multi-PDS model-checking problemfor PAs. Note that the final step of Alg. 2 can be performed with a single BDDoperation.

Theorem 3. For 2-PDS Π = (P1,P2, Locks, Σ) and PA A, there exists anexecution of Π that drives A to the accepting state iff Alg. 2 returns true.

Proof. See [13, App. D.2]. ut


input : A 2-PDS (P1,P2, Locks, Σ) and a PA A.output: true if there is an execution that drives A to the accepting state.let A1

post∗ ← post∗W1; let A2

post∗ ← post∗W2;1

let c1|Q|−1z|Q|−1 = A1

post∗`{〈(p1, q|Q|), u〉 | p1 ∈ P1 ∧ u ∈ Γ ∗1 }

´;2

let c2|Q|−1z|Q|−1 = A2

post∗`{〈(p2, q|Q|), u〉 | p2 ∈ P2 ∧ u ∈ Γ ∗2 }

´;3

return ∃〈LH0,cLH1〉 ∈ c1|Q|−1, 〈LH0,cLH2〉 ∈ c2|Q|−1 : Compatible(cLH1,cLH2);4

Algorithm 2: The symbolic decision procedure.

8 Experiments

Our experiment concerned detecting AS-serializability violations (or proving theirabsence) in models of concurrent Java programs. The experiment was designed tocompare the performance of Alg. 2 against that of the communicating-pushdownsystem (CPDS) semi-decision procedure from [10]. Alg. 2 was implemented usingthe Wali WPDS library [21] (the multi-arity relational weight domain is includedin the Wali release 3.0). The weight domain uses the BuDDy BDD library [22].All experiments were run on a dual-core 3 GHz Pentium Xeon processor with 4GB of memory.

We analyzed four Java programs from the ConTest benchmark suite [17]. Ourtool requires that the allocation site of interest be annotated in the source program.We annotated eleven of the twenty-seven programs that ConTest documentationidentifies as having “non-atomic” bugs. Our front-end currently handles eight ofthe eleven (the AST rewriting of [10] currently does not support certain Javaconstructs). Finally, after abstraction, four of the eight EML models did not uselocks, so we did not analyze them further. The four that we used in our studyare SoftwareVerificationHW, BugTester, BuggyProgram, and shop.

For each program, the front-end of the Empire tool [10] was used to createan EML program. An EML program has a set of shared-memory locations,SMem, a set of locks, SLocks, and a set of EML processes, SProcs. Five of thefourteen PAs used for detecting AS-serializability violations check behaviors thatinvolve a single shared-memory location; the other nine check behaviors thatinvolve a pair of shared-memory locations. For each of the five PAs that involvea single shared location, we ran one query for each m ∈ SMem. For each of thenine PAs that involve a pair of shared locations, we ran one query for each(m1,m2) ∈ SMem × SMem. In total, each tool ran 2,147 queries. Fig. 4 showslog-log scatter-plots of the execution times, classified into the 43 queries for whichAlg. 2 reported an AS-serializability violation (left-hand graph), and the 2,095queries for which Alg. 2 verified correctness (right-hand graph).

Although the CPDS-based method is a semi-decision procedure, it is capableof both (i) verifying correctness, and (ii) finding AS-serializability violations [10].(The third possibility is that it times out.) Comparing the total time to run allqueries, Alg. 2 ran 7.5X faster (136,235 seconds versus 17,728 seconds). The CPDS-based method ran faster than Alg. 2 on some queries, although never more thanabout 8X faster; in contrast, Alg. 2 was more than two orders of magnitude fasteron some queries. Moreover, the CPDS-based method timed out on about 68% of the


CPDS (seconds)

Alg

. 2 (

seconds)

2^−4

2^−2

2^0

2^2

2^4

2^6

2^8

2^−4 2^−2 2^0 2^2 2^4 2^6 2^8

AS−Violation

2^−4 2^−2 2^0 2^2 2^4 2^6 2^8

No AS−Violation

Fig. 4. Log-log scatter-plots of the execution times of Alg. 2 (y-axis) versus the CPDSsemi-decision procedure [10] (x-axis). The dashed lines denote equal running times;points below and to the right of the dashed lines are runs for which Alg. 2 was faster.The timeout threshold was 200 seconds; the minimal reported time is .25 seconds.The vertical bands near the right-hand axes represent queries for which the CPDSsemi-decision procedure timed out. (The horizontal banding is due to the fact that, fora given program, Alg. 2 often has similar performance for many queries.)

Query Category

CPDS succeeded CPDS timed outAlg. 2 succeeded Alg. 2 succeeded

Impl. (685 of 2,147) (1,453 of 2,147)

CPDS 6,006 130,229

Alg. 2 2,428 15,310

Fig. 5. Total time (in seconds) for exam-

ples classified according to whether CPDS

succeeded or timed out.

queries—both for the ones for whichAlg. 2 reported an AS-serializability vio-lation (29 timeouts out of 43 queries), aswell as the ones for which Alg. 2 verifiedcorrectness (1,425 timeouts out of 2,095queries). Alg. 2 exceeded the 200-secondtimeout threshold on nine queries. TheCPDS-based method also timed out onthose queries. When rerun with no time-out threshold, Alg. 2 solved each of thenine queries in 205–231 seconds.

Fig. 5 partitions the examples according to whether CPDS succeeded or timedout. The 1,453 examples on which CPDS timed out (col. 3 of Fig. 5) might besaid to represent “harder” examples. Alg. 2 required 15,310 seconds for these,which is about 3X more than the 1,453/685 × 2,428 = 5,150 seconds expectedif the queries in the two categories were of equal difficulty for Alg. 2. Roughlyspeaking, therefore, the data supports the conclusion that what is harder forCPDS is also harder for Alg. 2.


9 Related Work

The present paper introduces a different technique than that used by Kahlonand Gupta [7]. To decide the model-checking problem for PAs (as well as certaingeneralizations not discussed here), one needs to check pairwise reachability ofmultiple global configurations in succession. Our algorithm uses WPDS weightsthat are sets of lock-history tuples, whereas Kahlon and Gupta use sets of pairsof configuration automata.

There are similarities between the kind of splitting step needed by Qadeerand Rehof to enumerate states at a context switch [1] in context-boundedmodel checking and the splitting step on sets of automaton-pairs needed in thealgorithm of Kahlon and Gupta [7] to enumerate compatible configuration pairs[15]. Kahlon and Gupta’s algorithm performs a succession of pre∗ queries; aftereach one, it splits the resulting set of automaton-pairs to enforce the invariantthat succeeding queries are only applied to compatible configuration pairs. Incontrast, our algorithm (i) analyzes each PDS independently using one post∗

query per PDS, and then (ii) ties together the answers obtained from the differentPDSs by performing a single compatibility check on the sets of lock-history tuplesthat result. Because our algorithm does not need a splitting step on intermediateresults, it avoids enumerating compatible configuration pairs, thereby enablingBDD-based symbolic representations to be used throughout.

The Kahlon-Gupta decision procedure has not been implemented [15], soa direct performance comparison was not possible. It is left for future work todetermine whether our approach can be applied to the decidable sub-logics ofLTL identified in [7].

Our approach of using sets of tuples is similar in spirit to the use of matrix[2] and tuple [3] representations to address context-bounded model checking [1].In this paper, we bound the number of phases, but permit an unbounded numberof context switches and an unbounded number of lock acquisitions and releasesby each PDS. The decision procedure is able to explore the entire state space ofthe model; thus, our algorithm is able to verify properties of multi-PDSs insteadof just performing bug detection.

Dynamic pushdown networks (DPNs) [23] extend parallel PDSs with theability to create threads dynamically. Lammich et al. [24] present a generalizationof acquisition histories to DPNs with well-nested locks. Their algorithm useschained pre∗ queries, an explicit encoding of acquisition histories in the statespace, and is not implemented.

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