Fine-Grained Complexity of Safety Verification


We study the fine-grained complexity of Leader Contributor Reachability (\({\textsf {LCR}} \)) and Bounded-Stage Reachability (\({\textsf {BSR}} \)), two variants of the safety verification problem for shared memory concurrent programs. For both problems, the memory is a single variable over a finite data domain. Our contributions are new verification algorithms and lower bounds. The latter are based on the Exponential Time Hypothesis (\({\textsf {ETH}} \)), the problem \({\textsf {Set~Cover}} \), and cross-compositions. \({\textsf {LCR}} \) is the question whether a designated leader thread can reach an unsafe state when interacting with a certain number of equal contributor threads. We suggest two parameterizations: (1) By the size of the data domain \({\texttt {D}}\) and the size of the leader \({\texttt {L}}\), and (2) by the size of the contributors \({\texttt {C}}\). We present algorithms for both cases. The key techniques are compact witnesses and dynamic programming. The algorithms run in \({\mathcal {O}}^*(({\texttt {L}}\cdot ({\texttt {D}}+1))^{{\texttt {L}}\cdot {\texttt {D}}} \cdot {\texttt {D}}^{{\texttt {D}}})\) and \({\mathcal {O}}^*(2^{{\texttt {C}}})\) time, showing that both parameterizations are fixed-parameter tractable. We complement the upper bounds by (matching) lower bounds based on \({\textsf {ETH}} \) and \({\textsf {Set~Cover}} \). Moreover, we prove the absence of polynomial kernels. For \({\textsf {BSR}} \), we consider programs involving \({\texttt {t}}\) different threads. We restrict the analysis to computations where the write permission changes \({\texttt {s}}\) times between the threads. \({\textsf {BSR}} \) asks whether a given configuration is reachable via such an \({\texttt {s}}\)-stage computation. When parameterized by \({\texttt {P}}\), the maximum size of a thread, and \({\texttt {t}}\), the interesting observation is that the problem has a large number of difficult instances. Formally, we show that there is no polynomial kernel, no compression algorithm that reduces the size of the data domain \({\texttt {D}}\) or the number of stages \({\texttt {s}}\) to a polynomial dependence on \({\texttt {P}}\) and \({\texttt {t}}\). This indicates that symbolic methods may be harder to find for this problem.


We study the fine-grained complexity of two safety verification problems [1, 18, 32] for shared memory concurrent programs. The motivation to reconsider these problems are recent developments in fine-grained complexity theory [7, 12, 35, 39]. They suggest that classifications such as NP or even FPT are too coarse to explain the success of verification methods. Instead, it should be possible to identify the precise influence that parameters of the input have on the verification time. Our contribution confirms this idea. We give new verification algorithms for the two problems that, for the first time, can be proven optimal in the sense of fine-grained complexity theory. To state the results, we need some background. As we proceed, we explain the development of fine-grained complexity theory.

There is a well-known gap between the success that verification tools see in practice and the judgments about computational hardness that worst-case complexity is able to give. The applicability of verification tools steadily increases by tuning them towards industrial instances. The complexity estimation is stuck with considering the input size or at best assuming certain parameters to be constant. However, the latter approach is not very enlightening if the runtime is \(n^k\), where n is the input size and k the parameter.

The observation of a gap between practical algorithms and complexity theory is not unique to verification but made in every field that has to solve hard computational problems. Complexity theory has taken up the challenge to close the gap. So-called fixed-parameter tractability (FPT) [13, 15] proposes to identify parameters k so that the runtime is \(f(k) poly (n)\), where f is a computable function and \( poly(n) \) denotes any polynomial dependent on n. These parameters are powerful in the sense that they dominate the complexity.

For an FPT-result to be useful, function f should only be mildly exponential, and of course k should be small in the instances of interest. Intuitively, they are what one needs to optimize. Fine-grained complexity is the study of upper and lower bounds on the function. Indeed, the fine-grained complexity of a problem is written as \(O^*(f(k))\), emphasizing f and k and suppressing the polynomial part. For upper bounds, the approach is still to come up with an algorithm.

For lower bounds, fine-grained complexity has taken a new and very pragmatic perspective. For the problem of n-variable 3-SAT the best known algorithm runs in \({\mathcal {O}}(2^n)\), and this bound has not been improved since 1970. The idea is to take improvements on this problem as unlikely, known as the exponential-time hypothesis (ETH) [35]. Formally, it asserts that there is no \(2^{o(n)}\)-time algorithm for 3-SAT. ETH serves as a lower bound that is reduced to other problems [39]. An even stronger assumption about \(\mathsf {SAT}\), called SETH  [7, 35], and a similar one about Set Cover [12] allow for lower bounds like the absence of \({\mathcal {O}}^*((2-\delta )^n)\) algorithms.

In this work, we contribute fine-grained complexity results for verification problems on concurrent programs. The first problem is reachability for a leader thread that is interacting with an unbounded number of contributors (LCR) [18, 32]. We show that, assuming a parameterization by the size of the leader \({\texttt {L}}\) and the size of the data domain \({\texttt {D}}\), the problem can be solved in \({\mathcal {O}}^*(({\texttt {L}}\cdot ({\texttt {D}}+1))^{{\texttt {L}}\cdot {\texttt {D}}} \cdot {\texttt {D}}^{{\texttt {D}}})\). At the heart of the algorithm is a compression of computations into witnesses. To check reachability, our algorithm then iterates over candidates for witnesses and checks each of them for being a proper witness. Interestingly, we can formulate a variant of the algorithm that seems to be suited for large state spaces.

Using \({\textsf {ETH}} \), we show that the algorithm is (almost) optimal. Moreover, the problem is shown to have a large number of hard instances. Technically, there is no polynomial kernel [5, 6]. Experience with kernel lower bounds is still limited. This notion of hardness seems to indicate that symbolic methods are hard to apply. The lower bounds that we present share similarities with the reductions from [8, 29, 30].

If we consider the size \({\texttt {C}}\) of the contributors as a parameter, we obtain an \({\mathcal {O}}^*(2^{\texttt {C}})\) upper bound. Our algorithm is based on dynamic programming. We use the technique to solve a reachability problem on a graph that is shown to be a compressed representation for \({\textsf {LCR}} \). The compression is based on a saturation argument which is inspired by thread-modular reasoning [23, 24, 31, 34]. With the hardness assumption on \({\textsf {Set~Cover}} \) we show that the algorithm is indeed optimal. Moreover, we prove the absence of a polynomial kernel.

Parameterizations of \({\textsf {LCR}} \) involving just a single parameter \({\texttt {D}}\) or \({\texttt {L}}\) are intractable. We show that these problems are \({\textsf {W}} [1]\)-hard. This proves the existence of an \({\textsf {FPT}} \)-algorithm for those parameterizations unlikely.

The second problem we study generalizes bounded context switching. Bounded stage reachability (BSR) asks whether a state is reachable if there is a bound \({\texttt {s}}\) on the number of times the write permission is allowed to change between the threads [1]. Again, we show the new form of kernel lower bound. The result is tricky and highlights the power of the computation model.

The results are summarized by the table below. Main findings are highlighted in gray. We present two new algorithms for \({\textsf {LCR}} \). Moreover, we suggest kernel lower bounds as hardness indicators for verification problems. The corresponding lower bound for \({\textsf {BSR}} \) is particularly difficult to achieve.


The conference version of this paper appeared in [10]. The paper at hand presents new results. This includes an improved algorithm for \({\textsf {LCR}} \) running in \({\mathcal {O}}^*(2^{\texttt {C}})\) time instead of \({\mathcal {O}}^*(4^{\texttt {C}})\) and a new \((2-\delta )^{\texttt {C}}\) lower bound based on \({\textsf {Set~Cover}} \). Together, upper and lower bound show that the optimal algorithm for the problem has been found. Moreover, we give proofs for the intractability of certain parameterizations of \({\textsf {LCR}} \) and \({\textsf {BSR}} \). This justifies our choice of parameters.

We provide a full version of the paper at hand in [11], including missing proofs and details of formal constructions.

Related Work

Concurrent programs communicating through a shared memory and having a fixed number of threads have been extensively studied [2, 16, 27, 33]. The leader contributor reachability problem as considered in this paper was introduced as parametrized reachability in [32]. In [18], it was shown to be \({\textsf {NP}} \)-complete when only finite state programs are involved and \({\textsf {PSPACE}} \)-complete for recursive programs. In [36], the parameterized pairwise reachability problem was considered and shown to be decidable. Parameterized reachability under a variant of round robin scheduling was proven decidable in [38].

The bounded stage restriction on the computations of concurrent programs as considered here was introduced in [1]. The corresponding reachability problem was shown to be \({\textsf {NP}} \)-complete when only finite state programs are involved. The problem remains in \({\textsf {NEXP}} \)-time and \({\textsf {PSPACE}} \)-hard for a combination of counters and a single pushdown. The bounded stage restriction generalizes the concept of bounded context switching from [40], which was shown to be \({\textsf {NP}} \)-complete in that paper. In [9], \({\textsf {FPT}} \)-algorithms for bounded context switching were obtained under various parameterization. In [3], networks of pushdowns communicating through a shared memory were analyzed under topological restrictions.

There have been few efforts to obtain fixed-parameter tractable algorithms for automata and verification-related problems. FPT-algorithms for automata problems have been studied in [21, 22, 41]. In [14], model checking problems for synchronized executions on parallel components were considered and proven intractable. In [17], the notion of conflict serializability was introduced for the TSO memory model and an FPT-algorithm for checking serializability was provided. The complexity of predicting atomicity violation on concurrent systems was considered in [20]. The finding is that FPT-solutions are unlikely to exist. In [19], the problem of checking correctness of a program along a pattern is investigated. The authors conduct an analysis in several parameters. The results range from \({\textsf {NP}} \)-hardness even for fixed parameters to \({\textsf {FPT}} \)-algorithms.


We introduce our model for programs, which is fairly standard and taken from [1, 18, 32], and give the basics on fixed-parameter tractability.


A program consists of finitely many threads that access a shared memory. The memory is modeled to hold a single value at a time. Formally, a (shared memory) program is a tuple \({\mathcal {A}}= ( D , a^0, (P_i)_{i \in [1..t]})\). Here, \( D \) is the data domain of the memory and \(a^0\in D \) is the initial value. Threads are modeled as control-flow graphs that write values to or read values from the memory. These operations are captured by \({{ Op }}( D ) = \lbrace !a, ?a \mid a \in D \rbrace \). We use the notation \({{ W }}( D ) = \lbrace !a \mid a \in D \rbrace \) for the write operations and \({{ R }}( D ) = \lbrace ?a \mid a \in D \rbrace \) for the read operations. A thread \(P_{id}\) is a non-deterministic finite automaton \(({{ Op }}( D ), Q, q^0, \delta )\) over the alphabet of operations. The set of states is Q with \(q^0\in Q\) the initial state. The final states will depend on the verification task. The transition relation is \(\delta \subseteq Q \times ({{ Op }}( D ) \cup \{ \varepsilon \}) \times Q\). We extend it to words and also write \(q \xrightarrow {w} q'\) for \(q' \in \delta (q,w)\). Whenever we need to distinguish between different threads, we add indices and write \(Q_{id}\) or \(\delta _{id}\).

The semantics of a program is given in terms of labeled transitions between configurations. A configuration is a pair \(( pc , a) \in (Q_1 \times \dots \times Q_t) \times D \). The program counter \( pc \) is a vector that shows the current state \( pc (i)\in Q_i\) of each thread \(P_i\). Moreover, the configuration gives the current value in memory. We call \(\ c^0 = ( pc ^0, a^0)\) with \( pc ^0(i)=q^0_i\) for all \(i\in [1..t]\) the initial configuration. Let C denote the set of all configurations. The program’s transition relation among configurations \(\rightarrow \ \subseteq C\times ({{ Op }}( D )\cup \lbrace \varepsilon \rbrace )\times C\) is obtained by lifting the transition relations of the threads. To define it, let \( pc _1 = pc [i=q_i]\), meaning thread \(P_i\) is in state \(q_i\) and otherwise the program counter coincides with \( pc \). Let \( pc _2= pc [i=q_i']\). If thread \(P_i\) tries to read with the transition \(q_i \xrightarrow {?a} q_i'\), then \(( pc _1, a)\xrightarrow {?a}( pc _2, a)\). Note that the memory is required to hold the desired value. If the thread has the transition \(q_i\xrightarrow {!b} q_i'\), then \(( pc _1, a)\xrightarrow {!b}( pc _2, b)\). Finally, \(q_i\xrightarrow {\varepsilon }q_i'\) yields \(( pc _1, a)\xrightarrow {\varepsilon }( pc _2, a)\). The program’s transition relation is generalized to words, \(c \xrightarrow {w} c'\). We call such a sequence of consecutive labeled transitions a computation. To indicate that there is a word justifying a computation from c to \(c'\), we write \(c \rightarrow ^* c'\). We may use an index \(\xrightarrow {w}_i\) to indicate that the computation was induced by thread \(P_i\). Where appropriate, we also use the program as an index, \(\xrightarrow {w}_{\mathcal {A}}\).

Fixed-Parameter Tractability

We wish to study the fine-grained complexity of safety verification problems for the above programs. This means our goal is to identify parameters of these problems that satisfy two properties. First, in practical instances they are small. Second, assuming that these parameters are small, show that efficient verification algorithms can be obtained. Parametrized complexity is a branch of complexity theory that makes precise the idea of being efficient relative to a parameter.

Fix a finite alphabet \(\varSigma \). A parameterized problem L is a subset of \(\varSigma ^*\times {{\mathbb {N}}}\). The problem is fixed-parameter tractable if there is a deterministic algorithm that, given \((x,k)\in \varSigma ^*\times {{\mathbb {N}}}\), decides \((x,k)\in L\) in time \(f(k)\cdot |x|^{O(1)}\). We use \({\textsf {FPT}} \) for the class of all such problems and say a problem is \({\textsf {FPT}} \) to mean it is in that class. Note that f is a computable function only depending on the parameter k. It is common to denote the runtime by \({\mathcal {O}}^*(f(k))\) and suppress the polynomial part. We will be interested in the precise dependence on the parameter, in upper and lower bounds on the function f. This study is often referred to as fine-grained complexity.

Lower bounds on f are usually obtained from assumptions about \({\mathsf {SAT}}\). The most famous is the Exponential Time Hypothesis (\({\textsf {ETH}} \)). It assumes that there is no algorithm solving n-variable \({\textsf {3}} \text {-}{\textsf {SAT}} \) in \(2^{o(n)}\) time. Then, the reasoning is as follows: If f drops below a certain bound, \({\textsf {ETH}} \) would fail. Other standard assumptions for lower bounds are the Strong Exponential Time Hypothesis (\({\textsf {SETH}} \)) and the hardness assumption of \({\textsf {Set~Cover}} \). We postpone the definition of the latter and focus on \({\textsf {SETH}} \). This assumption is more restrictive than \({\textsf {ETH}} \). It asserts that n-variable \({\mathsf {SAT}}\) cannot be solved in \({\mathcal {O}}^*((2-\delta )^n)\) time for any \(\delta > 0\).

While many parameterizations of \({\textsf {NP}} \)-hard problems were proven to be fixed-parameter tractable, there are problems that are unlikely to be \({\textsf {FPT}} \). Such problems are hard for the complexity class \({\textsf {W}} [1]\). For a theory of relative hardness, the appropriate notion of reduction is called parameterized reduction. Given parameterized problems \(L, L'\subseteq \varSigma ^*\times {{\mathbb {N}}}\), we say that L is reducible to \(L'\) via a parameterized reduction if there is an algorithm that transforms an input (xk) to an input \((x',k')\) in time \(g(k) \cdot |x|^{O(1)}\) such that \((x,k) \in L\) if and only if \((x',k') \in L'\). Here, g is a computable function and \(k'\) is computed by a function only dependent on k.

Leader Contributor Reachability

We consider the leader contributor reachability problem for shared memory programs. The problem was introduced in [32] and shown to be \({\textsf {NP}} \)-complete in [18] for the finite state case.Footnote 1 We contribute two new verification algorithms that target two parameterizations of the problem. In both cases, our algorithms establish fixed-parameter tractability. Moreover, with matching lower bounds we prove them to be optimal even in the fine-grained sense.

An instance of the leader contributor reachability problem is given by a shared memory program of the form \({\mathcal {A}}= ( D , a^0, (P_L, (P_i)_{i \in [1..t]}))\). The program has a designated leader thread \(P_L\) and several contributor threads \(P_1, \dots , P_t\). In addition, we are given a set of unsafe states for the leader. The task is to check whether the leader can reach an unsafe state when interacting with a number of instances of the contributors. It is worth noting that the problem can be reduced to having a single contributor. Let the corresponding thread \(P_C\) be the union of \(P_1, \dots , P_t\) (constructed using an initial \(\varepsilon \)-transition). We base our complexity analysis on this simplified formulation of the problem.

For the definition, let \({\mathcal {A}}= ( D , a^0, (P_L,P_C))\) be a program with two threads. Let \(F_L \subseteq Q_L\) be a set of unsafe states of the leader. For any \(t \in {{\mathbb {N}}}\), define the program \({\mathcal {A}}^t = ( D , a^0, (P_L, (P_C)_{i\in [1..t]} ))\) to have exactly t copies of \(P_C\). Further, let \(C^f\) be the set of configurations where the leader is in an unsafe state (from \(F_L\)). The problem of interest is as follows:


We consider two parameterizations of \({\textsf {LCR}} \). First, we parameterize by \({\texttt {D}}\), the size of the data domain \( D \), and \({\texttt {L}}\), the number of states of the leader \(P_L\). We denote the parameterization by \({\textsf {LCR}} ({\texttt {D}}, {\texttt {L}})\). The second parameterization that we consider is \({\textsf {LCR}} ({\texttt {C}})\), a parameterization by the number of states of the contributor \(P_C\). For both, \({\textsf {LCR}} ({\texttt {D}},{\texttt {L}})\) and \({\textsf {LCR}} ({\texttt {C}})\), we present fine-grained analyses that include \({\textsf {FPT}} \)-algorithms as well as lower bounds for runtimes and kernels.

While for \({\textsf {LCR}} ({\texttt {D}}, {\texttt {L}})\) we obtain an \({\textsf {FPT}} \)-algorithm, it is not likely that \({\textsf {LCR}} ({\texttt {D}})\) and \({\textsf {LCR}} ({\texttt {L}})\) admit the same. We prove that these parameterizations are \({\textsf {W}} [1]\)-hard.

Parameterization by Memory and Leader

We give an algorithm that solves \({\textsf {LCR}} \) in time \({\mathcal {O}}^*( ({\texttt {L}}\cdot ({\texttt {D}}+ 1))^{{\texttt {L}}\cdot {\texttt {D}}} \cdot {\texttt {D}}^{{\texttt {D}}} )\), which means \({\textsf {LCR}} ({\texttt {D}}, {\texttt {L}})\) is \({\textsf {FPT}} \). We then show how to modify the algorithm to solve instances of \({\textsf {LCR}} \) as they are likely to occur in practice. Interestingly, the modified version of the algorithm lends itself to an efficient implementation based on off-the-shelf sequential model checkers. We conclude with lower bounds for \({\textsf {LCR}} ({\texttt {D}}, {\texttt {L}})\).

Upper Bound

We give an algorithm for the parameterization \({\textsf {LCR}} ({\texttt {D}}, {\texttt {L}})\). The key idea is to compactly represent computations that may be present in an instance of the given program. To this end, we introduce a domain of so-called witness candidates. The main technical result, Lemma 6, links computations and witness candidates. It shows that reachability of an unsafe state holds in an instance of the program if and only if there is a witness candidate that is valid (in a precise sense). With this, our algorithm iterates over all witness candidates and checks each of them for being valid. To state the overall result, let \({ Wit }({\texttt {L}}, {\texttt {D}}\,) = \left( {\texttt {L}}\cdot ({\texttt {D}}+ 1) \right) ^{{\texttt {L}}\cdot {\texttt {D}}} \cdot {\texttt {D}}^{{\texttt {D}}} \cdot {\texttt {L}}\) be the number of witness candidates and let \({ Valid }({\texttt {L}}, {\texttt {D}}, {\texttt {C}}\,) = {\texttt {L}}^3 \cdot {\texttt {D}}^{\, 2} \cdot {\texttt {C}}^2\) be the time it takes to check validity of a candidate. Note that it is polynomial.

Theorem 1

\({\textsf {LCR}} \) can be solved in time \({\mathcal {O}}( { Wit }({\texttt {L}}, {\texttt {D}}\,) \cdot { Valid }({\texttt {L}}, {\texttt {D}}, {\texttt {C}}\,) )\).

Let \({\mathcal {A}}= ( D , a^0, (P_L, P_C))\) be the program of interest and \(F_L\) be the set of unsafe states in the leader. Assume we are given a computation \(\rho \) showing that \(P_L\) can reach a state in \(F_L\) when interacting with a number of contributors. We explain the main ideas to find an efficient representation for \(\rho \) that still allows for the reconstruction of a similar computation. To simplify the presentation, we assume the leader never writes !a and immediately reads ?a (same value). If this is the case, the read can be replaced by \(\varepsilon \).

In a first step, we delete most of the moves in \(\rho \) that were carried out by the contributors. We only keep first writes. For each value a, this is the write transitions \( fw (a) = c \xrightarrow {!a} c'\) where a is written by a contributor for the first time. The reason we can omit subsequent writes of a is the following: If \( fw (a)\) is carried out by contributor \(P_1\), we can assume that there is an arbitrary number of other contributors that all mimicked the behavior of \(P_1\). This means whenever \(P_1\) did a transition, they copycatted it right away. Hence, there are arbitrarily many contributors pending to write a. Phrased differently, the symbol a is available for the leader whenever \(P_L\) needs to read it. The idea goes back to the Copycat Lemma stated in [18]. The reads of the contributors are omitted as well. We will make sure they can be served by the first writes and the moves done by \(P_L\).

After the deletion, we are left with a shorter expression \(\rho '\). We turn it into a word w over the alphabet \(Q_L \cup D _{\bot }\cup {\bar{ D }}\) with \( D _{\bot }= D \cup \lbrace \bot \rbrace \) and \({\bar{ D }}= \lbrace {\bar{a}} \mid a \in D \rbrace \). Each transition \(c \xrightarrow {!a / ?a / \varepsilon }_L c'\) in \(\rho '\) that is due to the leader moving from q to \(q'\) is mapped (i) to \(q.a.q'\) if it is a write and (ii) to \(q.\bot .q'\) otherwise. A first write \( fw (a) = c \xrightarrow {a} c'\) of a contributor is mapped to \({\bar{a}}\). We may assume that the resulting word w is of the form \(w = w_1 . w_2\) with \(w_1 \in ( ( Q_L . D _{\bot })^* . {\bar{ D }})^*\) and \(w_2 \in (Q_L . D _{\bot })^* . F_L\). Note that w can still be of unbounded length.

In order to find a witness of bounded length, we compress \(w_1\) and \(w_2\) to \(w'_1\) and \(w'_2\). Between two first writes \({\bar{a}}\) and \({\bar{b}}\) in \(w_1\), the leader can perform an unbounded number of transitions, represented by a word in \(( Q_L . D _{\bot })^*\). Hence, there are states \(q \in Q_L\) repeating between \({\bar{a}}\) and \({\bar{b}}\). We contract the word between the first and the last occurrence of q into just a single state q. This state now represents a loop on \(P_L\). Since there are \({\texttt {L}}\) states in the leader, this bounds the number of contractions. Furthermore, we know that the number of first writes is bounded by \({\texttt {D}}\), each symbol can be written for the first time at most once. Thus, the compressed string \(w'_1\) is a word in the language \(( (Q_L. D _{\bot })^{\le {\texttt {L}}} . {\bar{ D }})^{\le {\texttt {D}}}\).

The word \(w_2\) is of the form \(w_2 = q.u\) for a state \(q \in Q_L\) and a word u. We truncate the word u and only keep the state q. Then we know that there is a computation leading from q to a state in \(F_L\) where \(P_L\) can potentially write any symbol but read only those symbols which occurred as a first write in \(w'_1\). Altogether, we are left with a word of bounded length.

Definition 2

The set of witness candidates is \({\mathcal {E}}= ( (Q_L. D _{\bot })^{\le {\texttt {L}}} . {\bar{ D }})^{\le {\texttt {D}}} . Q_L\).

Before we elaborate on the precise relation between witness candidates and computations, we turn to an example. It shows how an actual computation is compressed to a witness candidate following the above steps.

Example 3

Consider the program \({\mathcal {A}}= ( D , a^0, (P_L,P_C))\) with domain \( D \), leader thread \(P_L\), and contributor thread \(P_C\) given in Fig. 1. We follow a computation in \({\mathcal {A}}^2\) that reaches the unsafe state \(q_4\) of the leader. Note that the transitions are labeled by L and C, depending on whether the leader or a contributor moved.

$$\begin{aligned} (q_0,p_0,p_0,a^0) \xrightarrow {{!a}}_C (q_0,p_1,p_0,a)&\xrightarrow {?a}_L (q_1,p_1,p_0,a) \xrightarrow {!b}_L \\ (q_2,p_1,p_0,b) \xrightarrow {?b}_C (q_2,p_1,p_2,b)&\xrightarrow {{!c}}_C (q_2,p_1,p_2,c) \xrightarrow {?c}_L \\ (q_3,p_1,p_2,c) \xrightarrow {!a}_C (q_3,p_1,p_2,a)&\xrightarrow {?a}_L (q_4,p_1,p_2,a). \end{aligned}$$

We construct a witness candidate out of the computation. To this end, we only keep the first writes of the contributors. These are the write !a in the first transition and the write !c in the fifth transition. Both are marked red. They will be represented in the witness candidate by the symbols \({\bar{a}}, {\bar{c}} \in {\bar{ D }}\).

Now we map the transitions of the leader to words. Writes are preserved, reads are mapped to \(\bot \). Then we obtain the witness candidate

$$\begin{aligned} {\bar{a}} \, . \, q_0 \, . \, \bot \, . \, q_1 \, . \, b \, . \, {\bar{c}} \, . \, q_2 . \end{aligned}$$

Note that we omit the last two transitions of the leader. The reason is as follows. After the first write \({\bar{c}}\), the leader is in state \(q_2\). From this state, the leader can reach \(q_4\) while only reading from first writes that have already appeared in the witness candidate, namely a and c. Hence, we can truncate the witness candidate at that point and do not have to keep the remaining computation to \(q_4\).

Fig. 1

Leader thread \(P_L\) (left) and contributor thread \(P_C\) (right) over the data domain \( D = \lbrace a^0, a, b, c \rbrace \). The only unsafe state of the leader is given by \(F_L = \lbrace q_4 \rbrace \)

To characterize computations in terms of witness candidates, we define the notion of validity. This needs some notation. Consider a word \(w = w_1 \dots w_\ell \) over some alphabet \(\varGamma \). For \(i \in [1..\ell ]\), we set \(w[i] = w_i\) and \(w[1..i] = w_1 \dots w_i\). If \(\varGamma ' \subseteq \varGamma \), we use \(w \! \downarrow _{\varGamma '}\) for the projection of w to the letters in \(\varGamma '\).

Consider a witness candidate \(w \in {\mathcal {E}}\) and let \(i \in [1..|w|]\). We use \({\bar{ D }}(w, i)\) for the set of all first writes that occurred in w up to position i. Formally, we define it to be \({\bar{ D }}(w,i) = \lbrace a \mid {\bar{a}} \text { is a letter in } w[1..i] \! \downarrow _{{\bar{ D }}} \rbrace \). We abbreviate \({\bar{ D }}(w, |w|)\) as \({\bar{ D }}(w)\). Let \(q \in Q_L\) and \(S \subseteq D \). Recall that the state represents a loop in \(P_L\). The set of all letters written within a loop from q to q when reading only symbols from S is \({\texttt {Loop}}(q,S) = \lbrace a \mid a \in D \text { and } \exists v_1, v_2 \in ({{ W }}( D ) \cup {{ R }}(S))^* : q \xrightarrow {v_1 !a v_2}_L q \rbrace \).

The definition of validity is given next. Technical details of the three requirements are made precise in the text below.

Definition 4

A witness candidate \(w \in {\mathcal {E}}\) is valid if it satisfies the following properties: (1) First writes are unique. (2) The word w encodes a run on \(P_L\). (3) There are supportive computations on the contributors.

  1. (1)

    If \(w \! \downarrow _{{\bar{ D }}} \, = \bar{c_1} \dots \bar{c_\ell }\), then the \(\bar{c_i}\) are pairwise different.

  2. (2)

    Let \(w \! \downarrow _{Q_L \cup D _{\bot }} = q_1 a_1 q_2 a_2 \dots a_\ell q_{\ell +1}\). If \(a_i \in D \), then \(q_i \xrightarrow {!a_i}_L q_{i+1} \in \delta _L\) is a write transition of \(P_L\). If \(a_i = \bot \), then we have an \(\varepsilon \)-transition \(q_i \xrightarrow {\varepsilon }_L q_{i+1}\). Alternatively, there is a read \(q_i \xrightarrow {?a}_L q_{i+1}\) of a symbol \(a \in {\bar{ D }}(w, {\texttt {pos}}(a_i))\) that already occurred within a first write (the leader does not read its own writes). Here, we use \({\texttt {pos}}(a_i)\) to access the position of \(a_i\) in w. State \(q_1=q^0_L\) is initial. There is a run from \(q_{\ell + 1}\) to a state \(q_f \in F_L\). During this run, reading is restricted to symbols that occurred as first writes in w. Formally, there is a word \(v \in ({{ W }}( D ) \cup {{ R }}({\bar{ D }}(w)))^*\) leading to an unsafe state \(q_f\). We have \(q_{\ell +1} \xrightarrow {v}_L q_f\).

  3. (3)

    For each prefix \(v{\bar{a}}\) of w with \({\bar{a}} \in {\bar{ D }}\) there is a computation \(q^0_C \xrightarrow {u !a}_C q\) on \(P_C\) so that the reads in u can be obtained from v. Formally, let \(u' = u \! \downarrow _{{{ R }}( D )}\). Then there is an embedding of \(u'\) into v, a monotone map \(\mu : [1..|u'|] \rightarrow [1..|v|]\) that satisfies the following. Let \(u'[i] =\ ?a\) with \(a \in D \). The read is served in one of the following three ways. We may have \(v[\mu (i)] = a\), which corresponds to a write of a by \(P_L\). Alternatively, \(v[\mu (i)] = q \in Q_L\) and \(a \in {\texttt {Loop}}(q,{\bar{ D }}(w,\mu (i)))\). This amounts to reading from a leader’s write that was executed in a loop. Finally, we may have \(a \in {\bar{ D }}(w, \mu (i))\), corresponding to reading from another contributor.

Our goal is to prove that a valid witness candidate exists if and only if there is a computation leading to an unsafe state. Before we state the corresponding lemma, we provide some intuition for the three requirements along an example.

Example 5

Reconsider the program \({\mathcal {A}}\) from Fig. 1. We elaborate on why the three requirements for validity are essential. To this end, we present three witness candidates, each violating one of the requirements. Thus, these candidates cannot correspond to an actual computation of the program.

The witness candidate \(w_1 = {\bar{a}} \, . \, q_0 \, . \, \bot \, . \, q_1 \, . \, b \, . \, {\bar{a}} \, . \, q_2\) clearly violates requirement (1) due to the repetition of \({\bar{a}}\). Since first writes are unique there cannot exist a computation of program \({\mathcal {A}}\) following candidate \(w_1\).

Requirement (2) asks for a proper run on the leader thread \(P_L\). Hence, the witness candidate \(w_2 = {\bar{a}} \, . \, q_0 \, . \, a \, . \, q_1 \, . \, b \, . \, {\bar{c}} \, . \, q_2\) violates the requirement although it satisfies (1). The subword \(q_0 \, . \, a \, . \, q_1\) of \(w_2\) encodes that the leader should take the transition \(q_0 \xrightarrow {!a}_L q_1\). But this transition does not exist in \(P_L\). Consequently, there is no computation of \({\mathcal {A}}\) which corresponds to the witness candidate \(w_2\).

For requirement (3), consider the candidate \(w_3 = {\bar{a}} \, . \, q_0 \, . \, \bot \, . \,q_1 \, . \, \bot \, . \, {\bar{c}} \, . \, q_2\). It clearly satisfies (1). Requirement (2) is also fulfilled. In fact, the subwords encoding transitions of the leader are \(q_0 \, . \, \bot \, . \, q_1\) and \(q_1 \, . \, \bot \, . \, q_2\). The first subword corresponds to transition \(q_0 \xrightarrow {?a}_L q_1\) which can be taken since a already appeared as a first write in \(w_3\). The second subword refers to the transition \(q_1 \xrightarrow {\varepsilon }_L q_2\).

To explain that \(w_3\) does not satisfy requirement (3), we show that c cannot be provided as a first write. To this end, assume that \(w_3\) satisfies (3). Then, for the prefix \(v . {\bar{c}}\) with \(v = {\bar{a}} \, . \, q_0 \, . \, \bot \, . \,q_1 \, . \, \bot \), there is a computation of the form \(p_0 \xrightarrow {u!c}_C p_2\). The reads in u are either first writes in v or writes provided by the leader (potentially in loops). Symbol b is not provided as such: It is neither a first write in v nor a symbol written by the leader (in a loop) along v. However, a computation u leading to state \(p_2\) in \(P_C\) needs to read b once. Hence, such a computation does not exist and c cannot be provided as a first write.

The witness candidate \(w = {\bar{a}} \, . \, q_0 \, . \, \bot \, . \, q_1 \, . \, b \, . \, {\bar{c}} \, . \, q_2\) from Example 3 satisfies all the requirements. In particular (3) is fulfilled since b is written by the leader in the transition \(q_1 \xrightarrow {!b} q_2\). Hence, in this case, c can be provided as a first write.

Lemma 6

There is a \(t \in {{\mathbb {N}}}\) so that \(c^0\rightarrow ^*_{{\mathcal {A}}^t}c\) with \(c\in C^f\) if and only if there is a valid witness candidate \(w \in {\mathcal {E}}\).

Our algorithm iterates over all witness candidates \(w \in {\mathcal {E}}\) and tests whether w is valid. The number of candidates \({ Wit }({\texttt {L}}, {\texttt {D}})\) is \(\left( {\texttt {L}}\cdot ({\texttt {D}}+ 1) \right) ^{{\texttt {L}}\cdot {\texttt {D}}} \cdot {\texttt {D}}^{{\texttt {D}}} \cdot {\texttt {L}}\). This is due to the fact that we can force a witness candidate to have maximum length via inserting padding symbols. Hence, the number of candidates constitutes the first factor of the complexity estimation stated in Theorem 1. The polynomial factor \({ Valid }({\texttt {L}}, {\texttt {D}}, {\texttt {C}})\) is due to the following lemma.

Lemma 7

Validity of \(w \in {\mathcal {E}}\) can be checked in time \({\mathcal {O}}( {\texttt {L}}^3 \cdot {\texttt {D}}^{\, 2} \cdot {\texttt {C}}^{\, 2} \, )\).

Practical Algorithm

We improve the above algorithm so that it should work well on practical instances. The idea is to factorize the leader along its strongly connected components (SCCs), the number of which is assumed to be small in real programs. Technically, our improved algorithm works with valid SCC-witnesses. They symbolically represent SCCs rather than loops in the leader. To state the complexity, we first define the straight line depth, the number of SCCs the leader may visit during a computation. The definition needs a graph construction.

Let \({\mathcal {V}}\subseteq {\bar{ D }}^{\le {\texttt {D}}}\) contain only words that do not repeat letters. Let \(r = {\bar{c}}_1 \dots {\bar{c}}_\ell \in {\mathcal {V}}\) and \(i \in [0..\ell ]\). By \(P_L \! \downarrow _{i}\) we denote the automaton obtained from \(P_L\) by removing all transitions that read a value outside \(\lbrace c_1, \dots , c_i \rbrace \). Let \({\texttt {SCC}}(P_L \! \downarrow _{i})\) denote the set of all SCCs in this automaton. We construct the directed graph \(G(P_L,r)\) as follows. The vertices are the SCCs of all \(P_L \! \downarrow _{i}\), \(i \in [0..\ell ]\). There is an edge between \(S, S' \in {\texttt {SCC}}(P_L \! \downarrow _{i})\), if there are states \(q \in S, q' \in S'\) with \(q \xrightarrow {?a / !a / \varepsilon } q'\) in \(P_L \! \downarrow _{i}\). If \(S \in {\texttt {SCC}}(P_L \! \downarrow _{i-1})\) and \(S' \in {\texttt {SCC}}(P_L \! \downarrow _{i})\), we only get an edge if we can get from S to \(S'\) by reading \(c_i\). Note that the graph is acyclic.

The depth d(r) of \(P_L\) relative to r is the length of the longest path in \(G(P_L,r)\). The straight line depth is \({\texttt {d}}= \max \lbrace d(r) \mid r \in {\mathcal {V}} \rbrace \). The number of SCCs \({\texttt {s}}\) is the size of \({\texttt {SCC}}(P_L \! \downarrow _{0})\). With these values at hand, the number of SCC-witness candidates (the definition of which can be found in the full version of the paper) can be bounded by \({ Wit }_{ SCC }({\texttt {s}}, {\texttt {D}}, {\texttt {d}}) \le ({\texttt {s}}\cdot ({\texttt {D}}+ 1))^{{\texttt {d}}} \cdot {\texttt {D}}^{{\texttt {D}}} \cdot 2^{{\texttt {D}}+ {\texttt {d}}}\). The time needed to test whether a candidate is valid is \({ Valid }_{ SCC }({\texttt {L}}, {\texttt {D}}, {\texttt {C}}, {\texttt {d}}) = {\texttt {L}}^2 \cdot {\texttt {D}}\cdot {\texttt {C}}^2 \cdot {\texttt {d}}^2\).

Theorem 8

\({\textsf {LCR}} \) can be solved in time \({\mathcal {O}}({ Wit }_{ SCC }({\texttt {s}}, {\texttt {D}}, {\texttt {d}}) \cdot { Valid }_{ SCC }({\texttt {L}}, {\texttt {D}}, {\texttt {C}}, {\texttt {d}}))\).

For this algorithm, what matters is that the leader’s state space is strongly connected. The number of states has limited impact on the runtime.

Lower Bound

We prove that the algorithm from Theorem 1 is only a root-factor away from being optimal: A \(2^{o( \sqrt{{\texttt {L}}\cdot {\texttt {D}}} \cdot \log ( {\texttt {L}}\cdot {\texttt {D}}) )}\)-time algorithm for \({\textsf {LCR}} \) would contradict \({\textsf {ETH}} \). We achieve the lower bound by a reduction from \({\textsf {k}} \times {\textsf {k}} ~ {\textsf {Clique}} \), the problem of finding a clique of size k in a graph the vertices of which are elements of a \(k \times k\) matrix. Moreover, the clique has to contain one vertex from each row. Unless \({\textsf {ETH}} \) fails, the problem cannot be solved in time \(2^{o(k \cdot \log (k))}\) [39].

Technically, we construct from an instance (Gk) of \({\textsf {k}} \times {\textsf {k}} ~ {\textsf {Clique}} \) an instance \(({\mathcal {A}}= ( D , a^0, (P_L, P_C)), F_L)\) of \({\textsf {LCR}} \) such that \({\texttt {D}}= {\mathcal {O}}(k)\) and \({\texttt {L}}= {\mathcal {O}}(k)\). Furthermore, we show that G contains the desired clique of size k if and only if there is a \(t \in {{\mathbb {N}}}\) such that \(c^0 \rightarrow ^*_{{\mathcal {A}}^t} c\) with \(c \in C^f\). Suppose we had an algorithm for \({\textsf {LCR}} \) running in time \(2^{o( \sqrt{{\texttt {L}}\cdot {\texttt {D}}} \cdot \log ( {\texttt {L}}\cdot {\texttt {D}}) )}\). Combined with the reduction, this would yield an algorithm for \({\textsf {k}} \times {\textsf {k}} ~ {\textsf {Clique}} \) with runtime \(2^{o(\sqrt{k^2} \cdot \log (k^2))} = 2^{o(k \cdot \log k)}\). But unless the exponential time hypothesis fails, such an algorithm cannot exist.

Proposition 9

\({\textsf {LCR}} \) cannot be solved in time \(2^{o( \sqrt{{\texttt {L}}\cdot {\texttt {D}}} \cdot \log ( {\texttt {L}}\cdot {\texttt {D}}) )}\) unless \({\textsf {ETH}} \) fails.

We assume that the vertices V of G are given by tuples (ij) with \(i,j \in [1..k]\), where i denotes the row and j denotes the column in the matrix. In the reduction, we need the leader and the contributors to communicate on the vertices of G. However, we cannot store tuples (ij) in the memory as this would cause a quadratic blow-up \({\texttt {D}}= {\mathcal {O}}(k^2)\). Instead, we communicate a vertex (ij) as a string \({\texttt {row}}(i).{\texttt {col}}(j)\). We distinguish between row- and column-symbols to avoid stuttering, the repeated reading of the same symbol. With this, it cannot happen that a thread reads a row-symbol twice and takes it for a column.

The program starts its computation with each contributor choosing a vertex (ij) to store. For simplicity, we denote a contributor storing the vertex (ij) by \(P_{(i,j)}\). Note that there can be copies of \(P_{(i,j)}\).

Since there are arbitrarily many contributors, the chosen vertices are only a superset of the clique we want to find. To cut away the false vertices, the leader \(P_L\) guesses for each row the vertex belonging to the clique. Contributors storing other vertices than the guessed ones will be switched off bit by bit. To this end, the program performs for each \(i \in [1..k]\) the following steps: If \((i,j_i)\) is the vertex of interest, \(P_L\) first writes \({\texttt {row}}(i)\) to the memory. Each contributor that is still active reads the symbol and moves on for one state. Then \(P_L\) communicates the column by writing \({\texttt {col}}(j_i)\). Again, the active contributors \(P_{(i',j')}\) read.

Upon transmitting \((i,j_i)\), the contributors react in one of the following three ways: (1) If \(i' \ne i\), the contributor \(P_{(i',j')}\) stores a vertex of a different row. The computation in \(P_{(i',j')}\) can only go on if \((i',j')\) is connected to \((i,j_i)\) in G. Otherwise it will stop. (2) If \(i' = i\) and \(j' = j_i\), then \(P_{(i',j')}\) stores exactly the vertex guessed by \(P_L\). In this case, \(P_{(i',j')}\) can continue its computation. (3) If \(i' = i\) and \(j' \ne j\), thread \(P_{(i',j')}\) stores a different vertex from row i. The contributor has to stop.

After k such rounds, there are only contributors left that store vertices guessed by \(P_L\). Furthermore, each two of these vertices are connected. Hence, they form a clique. To transmit this information to \(P_L\), each \(P_{(i,j_i)}\) writes \(\#_i\) to the memory, a special symbol for row i. After \(P_L\) has read the string \(\#_1 \dots \#_k\), it moves to its final state. A formal construction can be found in the full version of the paper.

Note that the size \({\mathcal {O}}(k)\) of the data domain cannot be avoided, even if we encoded the row and column symbols in binary. The reason is that \(P_L\) needs a confirmation of k contributors that were not stopped during the guessing and terminated correctly. Since contributors do not have final states, we need to transmit this information in the form of k different memory symbols.

Absence of a Polynomial Kernel

A kernelization of a parameterized problem is a compression algorithm. Given an instance, it returns an equivalent instance the size of which is bounded by a function only in the parameter. From an algorithmic perspective, kernels put a bound on the number of hard instances. Indeed, the search for small kernels is a key interest in algorithmics, similar to the search for \({\textsf {FPT}} \)-algorithms. It can be shown that kernels exist if and only if a problem admits an \({\textsf {FPT}} \)-algorithm [13].

Let Q be a parameterized problem. A kernelization of Q is an algorithm that given an instance (Bk), runs in polynomial time in B and k, and outputs an equivalent instance \((B', k')\) such that \(|B'| + k' \le g(k)\). Here, g is a computable function. If g is a polynomial, we say that Q admits a polynomial kernel.

Unfortunately, for many problems the community failed to come up with polynomial kernels. This lead to the contrary approach, namely disproving their existence [5, 6, 28]. The absence of a polynomial kernel constitutes an exponential lower bound on the number of hard instances. Like computational hardness results, such a bound is seen as an indication of general hardness of the problem. Technically, the existence of a polynomial kernel for the problem of interest is shown to imply \({\textsf {NP}} \subseteq {\textsf {co}} {\textsf {NP}} / {\mathsf {poly}}\). However, the inclusion is considered unlikely as it would cause a collapse of the polynomial hierarchy to the third level [42].

In order to link the existence of a polynomial kernel for \({\textsf {LCR}} ({\texttt {D}}, {\texttt {L}})\) with the above inclusion, we follow the framework developed in [6]. Let \(\varGamma \) be an alphabet. A polynomial equivalence relation is an equivalence relation \({{\mathcal {R}}}\) on \(\varGamma ^*\) with the following properties: Given \(x,y \in \varGamma ^*\), it can be decided in time polynomial in \(|x| + |y|\) whether \((x,y) \in {{\mathcal {R}}}\). Moreover, for \(n \in {{\mathbb {N}}}\) there are at most polynomially many equivalence classes in \({{\mathcal {R}}}\) restricted to \(\varGamma ^{\le n}\).

The key tool for proving kernel lower bounds are cross-compositions. Let \(L \subseteq \varGamma ^*\) be a language and \(Q \subseteq \varGamma ^* \times {\mathbb {N}}\) be a parameterized language. We say that L cross-composes into Q if there exists a polynomial equivalence relation \({{\mathcal {R}}}\) and an algorithm \({\mathcal {C}}\), together called the cross-composition, with the following properties: \({\mathcal {C}}\) takes as input \(\varphi _1, \dots , \varphi _I \in \varGamma ^*\), all equivalent under \({{\mathcal {R}}}\). It computes in time polynomial in \(\sum _{\ell =1}^{I} |\varphi _\ell |\) a string \((y,k) \in \varGamma ^* \times {{\mathbb {N}}}\) such that \((y,k) \in Q\) if and only if there is an \(\ell \in [1..I]\) with \(\varphi _\ell \in L\). Furthermore, parameter k is bounded by \(p(\max _{\ell \in [1..I]} |\varphi _\ell | + \log (I) )\), where p is a polynomial.

It was shown in [6] that a cross-composition of any \({\textsf {NP}} \)-hard language into a parameterized language Q prohibits the existence of a polynomial kernel for Q unless \({\textsf {NP}} \subseteq {\textsf {co}} {\textsf {NP}} / {\mathsf {poly}}\). In order to make use of this result, we show how to cross-compose \({\textsf {3}} \text {-}{\textsf {SAT}} \) into \({\textsf {LCR}} ({\texttt {D}}, {\texttt {L}})\). This yields the following:

Theorem 10

\({\textsf {LCR}} ({\texttt {D}}, {\texttt {L}})\) does not admit a polynomial kernel unless \({\textsf {NP}} \subseteq {\textsf {co}} {\textsf {NP}} / {\mathsf {poly}}\).

The difficulty in coming up with a cross-composition is the restriction on the size of the parameters. In our case, this affects \({\texttt {D}}\) and \({\texttt {L}}\): Both parameters are not allowed to depend polynomially on I, the number of given \({\textsf {3}} \text {-}{\textsf {SAT}} \)-instances. We resolve the polynomial dependence by encoding the choice of such an instance into the contributors via a binary tree.


(Idea) Assume some encoding of Boolean formulas as strings over a finite alphabet. We use the polynomial equivalence relation \({{\mathcal {R}}}\) defined as follows: Two strings \(\varphi \) and \(\psi \) are equivalent under \({{\mathcal {R}}}\) if both encode \({\textsf {3}} \text {-}{\textsf {SAT}} \)-instances, and the numbers of clauses and variables coincide.

Let the given \({\textsf {3}} \text {-}{\textsf {SAT}} \)-instances be \(\varphi _1, \dots , \varphi _I\). Every two of them are equivalent under \({{\mathcal {R}}}\). This means all \(\varphi _\ell \) have the same number of clauses m and use the same set of variables \(\lbrace x_1, \dots , x_n \rbrace \). We assume that \(\varphi _\ell = C^\ell _1 \wedge \dots \wedge C^\ell _m\).

We construct a program proceeding in three phases. First, it chooses an instance \(\varphi _\ell \), then it guesses an evaluation for all variables, and in the third phase it verifies that the evaluation satisfies \(\varphi _\ell \). While the second and the third phase do not cause a dependence of the parameters on I, the first phase does. It is not possible to guess a number \(\ell \in [1..I]\) and communicate it via the memory as this would provoke a polynomial dependence of \({\texttt {D}}\) on I.

To implement the first phase without a polynomial dependence, we transmit the indices of the \({\textsf {3}} \text {-}{\textsf {SAT}} \)-instances in binary. The leader guesses and writes tuples \((u_1, 1), \dots , (u_{\log (I)}, \log (I))\) with \(u_\ell \in \lbrace 0,1 \rbrace \) to the memory. This amounts to choosing an instance \(\varphi _\ell \) with binary representation \({\texttt {bin}}(\ell ) = u_1 \dots u_{\log (I)}\).

It is the contributors’ task to store this choice. Each time the leader writes a tuple \((u_i,i)\), the contributors read and branch either to the left, if \(u_i = 0\), or to the right, if \(u_i = 1\). Hence, in the first phase, the contributors are binary trees with I leaves, each leaf storing the index of an instance \(\varphi _\ell \). Since we did not assume that I is a power of 2, there may be computations arriving at leaves that do not represent proper indices. In this case, the computation deadlocks.

The size of \( D \) and \(P_L\) in the first phase is \({\mathcal {O}}(\log (I))\). Note that this satisfies the size-restrictions of a cross-composition.

For guessing the evaluation in the second phase, the program communicates on tuples \((x_i, v)\) with \(i \in [1..n]\) and \(v \in \lbrace 0,1 \rbrace \). The leader guesses such a tuple for each variable and writes it to the memory. Any participating contributor is free to read one of the tuples. After reading, it stores the variable and the evaluation.

In the third phase, the satisfiability check is performed as follows: Each contributor that is still active has stored in its current state the chosen instance \(\varphi _\ell \), a variable \(x_i\), and its evaluation \(v_i\). Assume that \(x_i\) when evaluated to \(v_i\) satisfies \(C^\ell _j\), the j-th clause of \(\varphi _\ell \). Then the contributor loops in its current state while writing the symbol \(\#_j\). The leader waits to read the string \(\#_1 \dots \#_m\). If \(P_L\) succeeds, we are sure that the m clauses of \(\varphi _\ell \) were satisfied by the chosen evaluation. Thus, \(\varphi _\ell \) is satisfiable and \(P_L\) moves to its final state. For details of the construction and a proof of correctness, we refer to the full version. \(\square \)

Parameterization by Contributors

The size of the contributors \({\texttt {C}}\) has substantial influence on the complexity of \({\textsf {LCR}} \). We show that the problem can be solved in time \({\mathcal {O}}^*(2^{\texttt {C}})\) via dynamic programming. Moreover, we present a matching lower bound proving it unlikely that \({\textsf {LCR}} \) can be solved in time \({\mathcal {O}}^*((2 - \delta )^{\texttt {C}})\), for any \(\delta > 0\). The result is obtained by a reduction from \({\textsf {Set~Cover}} \). Finally, a lower bound for the kernel of \({\textsf {LCR}} ({\texttt {C}})\) is provided.

Upper Bound

Our algorithm is based on dynamic programming. Intuitively, we cut a computation of the program along the states reached by the contributors. To this end, we keep a table with an entry for each subset of the contributors’ states. The entry of set \(S \subseteq Q_C\) contains those states of the leader that are reachable under a computation where the behavior of the contributors is limited to S. We fill the table by a dynamic programming procedure and check in the end whether a final state of the leader occurs in an entry. The result is as follows.

Theorem 11

\({\textsf {LCR}} \) can be solved in time \({\mathcal {O}}( 2^{\texttt {C}}\cdot {\texttt {C}}^{\, 4} \cdot {\texttt {L}}^2 \cdot {\texttt {D}}^{\, 2} )\).

To define the table, we first need a more compact way of representing computations that allows for fast iteration. The observation is that keeping one set of states for all contributors suffices. Let \(S \subseteq Q_C\) be the set of states reachable by the contributors in a given computation. By the Copycat Lemma [18], we can assume for each \(q \in S\) an arbitrary number of contributors that are currently in q. This means that we do not have to distinguish between different contributor instances.

Formally, we reduce the search space to \(V = Q_L \times D \times {{\mathcal {P}}}(Q_C)\). Instead of explicit configurations, we consider tuples (qaS), where \(q \in Q_L\), \(a \in D \), and \(S \subseteq Q_C\). Between these tuples, we define an edge relation E. If \(P_L\) writes \(a \in D \) with transition \(q \xrightarrow {!a} q'\), we get \((q, b, S) \rightarrow _E (q', a, S)\) for each \(b \in D \) and \(S \subseteq Q_C\). Reads of the leader are similar. Contributors also change the memory but saturate set S instead of changing the state: If there is a transition \(p \xrightarrow {!a} p'\) in \(P_C\) with \(p \in S\), then \((q, b, S) \rightarrow _E (q, a, S \cup \lbrace p' \rbrace )\) for each \(b \in D \) and \(q \in Q_L\). Reads are similar.

The set V together with the relation E form a finite directed graph \({\mathcal {G}}= (V,E)\). We call the node \(v^0= (q^0_L, a^0, \lbrace q^0_C \rbrace )\) the initial node. Computations are represented by paths in \({\mathcal {G}}\) starting in \(v^0\). Hence, we reduced \({\textsf {LCR}} \) to the problem of checking whether the set of nodes \(F_L \times D \times {{\mathcal {P}}}(Q_C)\) is reachable from the initial node in \({\mathcal {G}}\).

Lemma 12

There is a \(t \in {{\mathbb {N}}}\) so that \(c^0 \rightarrow ^*_{{\mathcal {A}}^t} c\) with \(c \in C^f\) if and only if there is a path in \({\mathcal {G}}\) from \(v^0\) to a node in \(F_L \times D \times {{\mathcal {P}}}(Q_C)\).

Before we elaborate on the algorithm solving reachability on \({\mathcal {G}}\) we turn to an example. It shows how \({\mathcal {G}}\) is constructed from a program and illustrates Lemma 12.

Fig. 2

Leader thread \(P_L\) (left) and contributor thread \(P_C\) (right). The data domain is given by \( D = \lbrace a^0, a, b, c \rbrace \) and the only unsafe state is \(F_L = \lbrace q_3 \rbrace \)

Example 13

We consider the program \({\mathcal {A}}= ( D , a^0, (P_L,P_C))\) from Fig. 2. The nodes of the corresponding graph \({\mathcal {G}}\) are given by \(V = Q_L \times D \times {{\mathcal {P}}}(\lbrace p_0,p_1,p_2 \rbrace )\). Its edges E are constructed following the above rules. For instance, we get an edge \((q_1,a,\lbrace p_0 \rbrace ) \rightarrow _E (q_1,a,\lbrace p_0,p_1 \rbrace )\) since \(P_C\) has a read transition \(p_0 \xrightarrow {?a} p_1\). Intuitively, the edge describes that currently, the leader is in state \(q_1\), the memory holds a, and an arbitrary number of contributors is waiting in \(p_0\). Then, some of these read a and move to \(p_1\). Hence, we might assume an arbitrary number of contributors in the states \(p_0\) and \(p_1\).

The complete graph \({\mathcal {G}}\) is presented in Fig. 3. For the purpose of readability, we only show the nodes that are reachable from \(v^0= (q_0,a^0,\lbrace p_0 \rbrace )\). Moreover, we omit self-loops and we present the graph as a collection of subgraphs. The latter means that for each subset S of \({{\mathcal {P}}}(\lbrace p_0,p_1,p_2 \rbrace )\), we consider the induced subgraph \({\mathcal {G}}[Q_L \times D \times \lbrace S \rbrace ]\). It contains the nodes \(Q_L \times D \times \lbrace S \rbrace \) and all edges that start and end in this set. Note that we omit the last component from a node (qaS) in \({\mathcal {G}}[Q_L \times D \times \lbrace S \rbrace ]\). The induced subgraphs are connected by edges that saturate S.

The red marked nodes are those which contain the unsafe state \(q_3\) of the leader. Consider a path from \(v^0\) to one of these nodes. It starts in \({\mathcal {G}}[Q_L \times D \times \lbrace \lbrace p_0 \rbrace \rbrace ]\). To reach one of the red nodes, the path has to traverse via \({\mathcal {G}}[Q_L \times D \times \lbrace \lbrace p_0,p_1 \rbrace \rbrace ]\) to \({\mathcal {G}}[Q_L \times D \times \lbrace Q_C \rbrace ]\) or via \({\mathcal {G}}[Q_L \times D \times \lbrace \lbrace p_0,p_2 \rbrace \rbrace ]\). Phrased differently, the states of the contributors need to be saturated two times along the path. This means that in an actual computation, there must be contributors in \(p_0\), \(p_1\), and \(p_2\). These can then provide the symbols b and c which are needed by the leader to reach \(q_3\).

Fig. 3

Graph \({\mathcal {G}}\) summarizing the computations of the program in Fig. 2. Self-loops and nodes not reachable from \(v^0= (q_0,a^0,\lbrace p_0 \rbrace )\) are omitted. We further omit the third component of nodes since it is clear from the context. Nodes that are marked red involve the unsafe state \(q_3\) of the leader. The blue highlighted area shows the slice \({\mathcal {G}}_{\lbrace p_0 \rbrace ,\lbrace p_0,p_1 \rbrace }\). (Color figure online)

Constructing \({\mathcal {G}}\) for a program and solving reachability takes time \({\mathcal {O}}^*(4^{\texttt {C}})\) [10]. Hence, we have to solve reachability without constructing \({\mathcal {G}}\) explicitly. Our algorithm computes a table T which admits a recurrence relation that simplifies the reachability query: Instead of solving reachability directly on \({\mathcal {G}}\), we can restrict to so-called slices of \({\mathcal {G}}\). These are subgraphs of polynomial size where reachability queries can be decided efficiently.

We define the table T. For each set \(S \subseteq Q_C\), we have an entry T[S], given by:

$$\begin{aligned} T[S] = \lbrace (q,a) \in Q_L \times D \mid v^0 \rightarrow ^*_E (q,a,S) \rbrace . \end{aligned}$$

Intuitively, T[S] contains all nodes in \({\mathcal {G}}[Q_L \times D \times \lbrace S \rbrace ]\) that are reachable from \(v^0\).

Assume we have already computed T. By Lemma 12 we get: There is a \(t \in {{\mathbb {N}}}\) so that \(c^0 \rightarrow ^*_{{\mathcal {A}}^t} c \in C^f\) if and only if there is an \(S \subseteq Q_C\) such that \(T[S] \cap F_L \times D \ne \emptyset \). The latter can be checked in time \({\mathcal {O}}(2^{\texttt {C}}\cdot {\texttt {L}}^2 \cdot {\texttt {D}}^2)\) as there are \(2^{\texttt {C}}\) candidates for S.

It remains to compute the table. Our goal is to employ a dynamic programming based on a recurrence relation over T. To formulate the relation, we need the notion of slices of \({\mathcal {G}}\). Let \(W \subseteq Q_C\) be a subset and \(p \in Q_C \setminus W\) be a state. We denote by S the union \(S = W \cup \lbrace p \rbrace \). The slice \({\mathcal {G}}_{W,S}\) is the induced subgraph \({\mathcal {G}}[ Q_L \times D \times \lbrace W,S \rbrace ]\). We denote its set of edges by \(E_{W,S}\).

The main idea of the recurrence relation is saturation. When traversing a path \(\pi \) in \({\mathcal {G}}\), the set of contributor states gets saturated over time. Assume we cut \(\pi \) each time after a new state gets added. Then we obtain subpaths, each being a path in a slice: If \(p \in Q_C\) gets added to \(W \subseteq Q_C\), the corresponding subpath is in \({\mathcal {G}}_{W,W\cup \lbrace p \rbrace }\). Phrased differently, for a set \(S \subseteq Q_C\), the entry T[S] is the union of those nodes that are reachable from \(T[S \setminus \lbrace p \rbrace ]\) in \({\mathcal {G}}_{S\setminus \lbrace p \rbrace ,S}\), for each \(p \in S\).

Formally, we define sets R(WS) for each \(W \subseteq Q_C\), \(p \in Q_C \setminus W\), and \(S = W \cup \lbrace p \rbrace \). These sets collect the nodes that are reachable from T[W] in the slice \({\mathcal {G}}_{W,S}\):

$$\begin{aligned} R(W,S) = \lbrace (q,a) \in Q_L \times D \mid \exists (q',a') \in T[W] \text { with } (q',a',W) \rightarrow ^*_{E_{W,S}} (q,a,S) \rbrace . \end{aligned}$$

Lemma 14

Table T admits the recurrence relation \(T[S] = \bigcup _{p \in S} R(S \setminus \lbrace p \rbrace ,S)\).

We illustrate the lemma and the introduced notions on an example. Afterwards, we show how to compute the table T by exploiting the recurrence relation.

Example 15

Reconsider the program given in Fig. 2. The table T has eight entries, one for each subset of \(Q_C\). The entries that are non-empty can be seen in the graph of Fig. 3. Each of the subgraphs contains exactly those nodes that are reachable from \(v^0\). For instance \(T[\lbrace p_0,p_1 \rbrace ] = \lbrace (q_1,a),(q_1,c) \rbrace \).

Let \(W = \lbrace p_0 \rbrace \) and \(S = \lbrace p_0,p_1 \rbrace \). Then, the slice \({\mathcal {G}}_{W,S}\) is shown in the figure as blue highlighted area. Note that it also contains the edge from \((q_1,a,\lbrace p_0 \rbrace )\) to \((q_1,a,\lbrace p_0,p_1 \rbrace )\), leading from \({\mathcal {G}}[Q_L \times D \times \lbrace \lbrace p_0 \rbrace \rbrace ]\) to \({\mathcal {G}}[Q_L \times D \times \lbrace \lbrace p_0,p_1 \rbrace \rbrace ]\).

The set R(WS) contains those nodes in \({\mathcal {G}}[Q_L \times D \times \lbrace S \rbrace ]\) that are reachable from T[W] in the slice \({\mathcal {G}}_{W,S}\). According to the graph, these are \((q_1,a,S)\) and \((q_1,c,S)\) and hence we get \(T[S] = R(W,S)\).

In general, not all nodes in T[S] are reachable from a single set T[W]. But if a node is reachable, then it is reachable from some set \(T[S\setminus \lbrace p \rbrace ]\) with \(p \in S\). Note that this is covered by the recurrence relation in Lemma 14. It branches over all sets \(S \setminus \lbrace p \rbrace \) and hence collects all nodes that are reachable from such a set.

We apply the recurrence relation in a bottom-up dynamic programming to fill the table T. Let \(S \subseteq Q_C\) be a subset and assume we already know \(T[S \setminus \lbrace p \rbrace ]\), for each \(p \in S\). Then, for a fixed p, we compute \(R(S \setminus \lbrace p \rbrace ,S)\) by a fixed-point iteration on the slice \({\mathcal {G}}_{S \setminus \lbrace p \rbrace ,S}\). The number of nodes in the slice is \({\mathcal {O}}({\texttt {L}}\cdot {\texttt {D}})\). Hence, the iteration takes time at most \({\mathcal {O}}({\texttt {L}}^2 \cdot {\texttt {D}}^2)\). It is left to construct \({\mathcal {G}}_{S \setminus \lbrace p \rbrace , S}\). We state the time needed in the following lemma. The proof is postponed so as to finish the complexity estimation of Theorem 11.

Lemma 16

Slice \({\mathcal {G}}_{S \setminus \lbrace p \rbrace , S}\) can be constructed in time \({\mathcal {O}}({\texttt {C}}^{\, 3} \cdot {\texttt {L}}^2 \cdot {\texttt {D}}^{\,2})\).

Wrapping up, we need \({\mathcal {O}}( {\texttt {C}}^3 \cdot {\texttt {L}}^2 \cdot {\texttt {D}}^2 )\) time for computing a set \(R(S \setminus \lbrace p \rbrace ,S)\). Due to the recurrence relation of Lemma 14, we have to compute at most \({\texttt {C}}\) sets \(R(S \setminus \lbrace p \rbrace ,S)\) for a given \(S \subseteq Q_C\). Hence, an entry T[S] can be computed in time \({\mathcal {O}}({\texttt {C}}^4 \cdot {\texttt {L}}^2 \cdot {\texttt {D}}^2)\). The estimation also covers the base case \(S = \lbrace p_C^0 \rbrace \), where T[S] can be computed by a fixed-point iteration on the induced subgraph \({\mathcal {G}}[Q_L \times D \times \lbrace S \rbrace ]\). Since the table T has \(2^{\texttt {C}}\) entries, the complexity estimation of Theorem 11 follows. It is left to prove Lemma 16.


The slice \({\mathcal {G}}_{S \setminus \lbrace p \rbrace ,S}\) consists of the two subgraphs \({\mathcal {G}}_{S \setminus \lbrace p \rbrace } = {\mathcal {G}}[Q_L \times D \times \lbrace S \setminus \lbrace p \rbrace \rbrace ]\) and \({\mathcal {G}}_{S} = {\mathcal {G}}[Q_L \times D \times \lbrace S \rbrace ]\), and the edges leading from \({\mathcal {G}}_{S \setminus \lbrace p \rbrace }\) to \({\mathcal {G}}_{S}\). We elaborate on how to construct \({\mathcal {G}}_{S}\). The construction of \({\mathcal {G}}_{S \setminus \lbrace p \rbrace }\) is similar.

First, we write down the nodes of \({\mathcal {G}}_{S}\). This can be done in time \({\mathcal {O}}({\texttt {L}}\cdot {\texttt {D}})\). Edges in the graph are either induced by transitions of the leader or by the contributor. The former ones can be added in time \({\mathcal {O}}(|\delta _L| \cdot {\texttt {D}}) = {\mathcal {O}}({\texttt {L}}^2 \cdot {\texttt {D}}^2)\) since a single transition of \(P_L\) may lead to \({\texttt {D}}\) edges. To add the latter edges, we browse \(\delta _C\) for transitions of the form \(s \xrightarrow {!a} s'\) with \(s,s' \in S\). Each such transition may induce \({\texttt {L}}\cdot {\texttt {D}}\) edges. Adding them takes time \({\mathcal {O}}(|\delta _C| \cdot {\texttt {C}}\cdot {\texttt {L}}\cdot {\texttt {D}}) = {\mathcal {O}}({\texttt {C}}^3 \cdot {\texttt {L}}\cdot {\texttt {D}}^2)\) since we have to test membership of \(s,s'\) in S. Note that we can omit transitions \(s \xrightarrow {?a} s'\) with \(s,s' \in S\) as their induced edges are self-loops in \({\mathcal {G}}_{S}\).

To complete the construction, we add the edges from \({\mathcal {G}}_{S \setminus \lbrace p \rbrace }\) to \({\mathcal {G}}_{S}\). These are induced by transitions \(r \xrightarrow {?a / !a} p \in \delta _C\) with \(r \in S \setminus \lbrace p \rbrace \). Since each of these may again lead to \({\texttt {L}}\cdot {\texttt {D}}\) different edges, adding all of them takes time \({\mathcal {O}}({\texttt {C}}^3 \cdot {\texttt {L}}\cdot {\texttt {D}}^2)\). In total, we estimate the time for the construction by \({\mathcal {O}}({\texttt {C}}^3 \cdot {\texttt {L}}^2 \cdot {\texttt {D}}^2)\). \(\square \)

Lower Bound

We prove it unlikely that \({\textsf {LCR}} \) can be solved in \({\mathcal {O}}^*((2-\delta )^{\texttt {C}})\) time, for any \(\delta > 0\). This shows that the algorithm from Sect. 3.2 has an optimal runtime. The lower bound is achieved by a reduction from \({\textsf {Set~Cover}} \), one of the 21 original \({\textsf {NP}} \)-complete problems by Karp [37]. We state its definition.


Besides its \({\textsf {NP}} \)-completeness, it is known that \({\textsf {Set~Cover}} \) admits an \({\mathcal {O}}^*(2^n)\)-time algorithm [25], where n is the size of the universe U. However, no algorithm solving \({\textsf {Set~Cover}} \) in time \({\mathcal {O}}^*((2-\delta )^n)\) for a \(\delta > 0\) is known so far. Actually, it is conjectured in [12] that such an algorithm cannot exist unless the \({\textsf {SETH}} \) breaks.

While a proof for the conjecture in [12] is still missing, the authors provide evidence in the form of relative hardness. They obtain lower bounds for prominent problems by tracing back to the assumed lower bound of \({\textsf {Set~Cover}} \). These bounds were not known before since \({\textsf {SETH}} \) is hard to apply: No suitable reductions from \({\mathsf {SAT}}\) to these problems are known so far. Hence, \({\textsf {Set~Cover}} \) can be seen as an alternative source for lower bounds whenever \({\textsf {SETH}} \) seems out of reach. This made the problem a standard assumption for hardness which is widely used [4, 9, 12].

To obtain the desired lower bound for \({\textsf {LCR}} \), we establish a polynomial time reduction from \({\textsf {Set~Cover}} \) that strictly preserves the parameter n. Formally, if \(({\mathcal {F}},U,r)\) is an instance of \({\textsf {Set~Cover}} \), we construct an instance \(({\mathcal {A}}= ( D , a^0, (P_L, P_C)), F_L)\) of \({\textsf {LCR}} \) where \({\texttt {C}}= n + c\) with c a constant. Note that even a linear dependence on n is not allowed. Moreover, the instance satisfies the equivalence: There is a set cover if and only if there is a \(t \in {{\mathbb {N}}}\) such that \(c^0 \rightarrow ^*_{{\mathcal {A}}^t} c\) with \(c \in C^f\). Assume we had an \({\mathcal {O}}^*((2-\delta )^{\texttt {C}})\)-time algorithm for \({\textsf {LCR}} \). With the reduction, this would immediately yield an \({\mathcal {O}}^*((2-\delta )^{n+c}) = {\mathcal {O}}^*((2-\delta )^n)\)-time algorithm for \({\textsf {Set~Cover}} \).

Proposition 17

If \({\textsf {LCR}} \) can be solved in \({\mathcal {O}}^*((2-\delta )^{\texttt {C}})\) time for a \(\delta > 0\), then \({\textsf {Set~Cover}} \) can be solved in \({\mathcal {O}}^*((2-\delta )^n)\) time.

For the proof of the proposition, we elaborate on the aforementioned reduction. The main idea is the following: We let the leader guess r sets from \({\mathcal {F}}\). The contributors store the elements that got covered by the chosen sets. In a final communication phase, the leader verifies that it has chosen a valid cover by querying whether all elements of U have been stored by the contributors.

Leader and contributors essentially communicate over the elements of U. For guessing r sets from \({\mathcal {F}}\), the automaton \(P_L\) consists of r similar phases. Each phase starts with \(P_L\) choosing an internal transition to a set \(S \in {\mathcal {F}}\). Once S is chosen, the leader writes a sequence of all \(u \in S\) to the memory.

A contributor in the program consists of \({\texttt {C}}= n+1\) states: An initial state and a state for each \(u \in U\). When \(P_L\) writes an element \(u \in S\) to the memory, there is a contributor storing this element in its states by reading u. Hence, each element that got covered by S is recorded in one of the contributors.

After r rounds of guessing, the contributors hold those elements of U that are covered by the chosen sets. Now the leader verifies that it has really picked a cover of U. To this end, it needs to check whether all elements of U have been stored by the contributors. Formally, the leader can only proceed to its final state if it can read the symbols \(u^\#\), for each \(u \in U\). A contributor can only write \(u^\#\) to the memory if it stored the element u before. Hence, \(P_L\) reaches its final state if and only if a valid cover of U was chosen.

Absence of a Polynomial Kernel

We prove that \({\textsf {3}} \text {-}{\textsf {SAT}} \) can be cross-composed into \({\textsf {LCR}} ({\texttt {C}})\). This shows that the problem is unlikely to admit a polynomial kernel. The result is the following.

Proposition 18

\({\textsf {LCR}} ({\texttt {C}}\, )\) does not admit a polynomial kernel unless \({\textsf {NP}} \subseteq {\textsf {co}} {\textsf {NP}} / {\mathsf {poly}}\).

For the cross-composition, let \(\varphi _1, \dots , \varphi _I\) be the given \({\textsf {3}} \text {-}{\textsf {SAT}} \)-instances, each two equivalent under \({{\mathcal {R}}}\), where \({{\mathcal {R}}}\) is the polynomial equivalence relation from Theorem 10. Then, each formula has the same number of clauses m and variables \(x_1, \dots , x_n\). Let us fix the notation to be \(\varphi _\ell = C^\ell _1 \wedge \dots \wedge C^\ell _m\).

The basic idea is the following. Leader \(P_L\) guesses the formula \(\varphi _\ell \) and an evaluation for the variables. The contributors store the latter. At the end, leader and contributors verify that the chosen evaluation indeed satisfies formula \(\varphi _\ell \).

For guessing \(\varphi _\ell \), the leader has a branch for each instance. Note that we can afford the size of the leader to depend on I since the cross-composition only restricts parameter \({\texttt {C}}\). Hence, we do not face the problem we had in Theorem 10.

Guessing the evaluation of the variables is similar to Theorem 10: The leader writes tuples \((x_i,v_i)\) with \(v_i \in \lbrace 0,1 \rbrace \) to the memory. The contributors store the evaluation in their states. After this guessing-phase, the contributors can write the symbols \(\#^\ell _j\), depending on whether the currently stored variable with its evaluation satisfies clause \(C^\ell _j\). As soon as the leader has read the complete string \(\#^\ell _1 \dots \#^\ell _m\), it moves to its final state, showing that the evaluation satisfied all clauses of \(\varphi _\ell \).

Note that parameter \({\texttt {C}}\) is of size \({\mathcal {O}}(n)\) and does not depend on I at all. Hence, the size-restrictions of a cross-composition are met.


We show the \({\textsf {W}} [1]\)-hardness of \({\textsf {LCR}} ({\texttt {D}})\) and \({\textsf {LCR}} ({\texttt {L}})\). Both proofs rely on a parameterized reduction from \({\textsf {k \text {-} Clique}} \), the problem of finding a clique of size k in a given graph. This problem is known to be \({\textsf {W}} [1]\)-complete [15]. Our result is the following.

Proposition 19

Both parameterizations, \({\textsf {LCR}} ({\texttt {D}})\) and \({\textsf {LCR}} ({\texttt {L}})\), are \({\textsf {W}} [1]\)-hard.

We first reduce \({\textsf {k \text {-} Clique}} \) to \({\textsf {LCR}} ({\texttt {L}})\). More precisely, we construct from an instance (Gk) of \({\textsf {k \text {-} Clique}} \) in polynomial time an instance \(({\mathcal {A}}= ( D , a^0, (P_L, P_C)), F_L)\) of \({\textsf {LCR}} \) with \({\texttt {L}}= {\mathcal {O}}(k)\). This meets the requirements of a parameterized reduction.

Program \({\mathcal {A}}\) operates in three phases. In the first phase, the leader chooses k vertices of the graph and writes them to the memory. Formally, it writes a sequence \((v_1,1). (v_2,2) \dots (v_k,k)\) with \(v_i\) vertices of G. During this selection, the contributors non-deterministically choose to store a suggested vertex \((v_i,i)\) in their state space.

In the second phase, the leader again writes a sequence of vertices using different symbols: \((w^\#_1,1) (w^\#_2,2) \dots (w^\#_k,k)\). Note that the vertices \(w_i\) do not have to coincide with the vertices from the first phase. It is then the contributor’s task to verify that the new sequence constitutes a clique. To this end, for each i, the program does the following: If a contributor storing \((v_i,i)\) reads the value \((w^\#_i,i)\), the computation on the contributor can only continue if \(w_i = v_i\). If a contributor storing \((v_j,j)\) with \(j \ne i\) reads \((w^\#_i,i)\), the computation can only continue if \(v_j \ne w_i\) and if there is an edge between \(v_j\) and \(w_i\).

Finally, in the third phase, we need to ensure that there was at least one contributor storing \((v_i,i)\) and that the above checks were all positive. To this end, a contributor that has successfully gone through the second phase and stores \((v_i,i)\) writes the symbol \(\#_i\) to the memory. The leader pends to read the sequence \(\#_1 \dots \#_k\). This ensures the selection of k different vertices, each two being adjacent.

For proving \({\textsf {W}} [1]\)-hardness of \({\textsf {LCR}} ({\texttt {D}})\), we reuse the above construction. However, the size of the data domain is \(|V| \cdot k\), where V is the set of vertices of G. Hence, it is not a parameterized reduction for parameter \({\texttt {D}}\). The factor \(|V|\) appears since leader and contributors communicate on the pure vertices. The main idea of the new reduction is to decrease the size of \( D \) by transmitting the vertices in binary. To this end, we add binary branching trees to the contributors that decode a binary encoding. We omit the details and refer to the full version of the paper.

Bounded-Stage Reachability

The bounded-stage reachability problem is a simultaneous reachability problem. It asks whether all threads of a program can reach an unsafe state when restricted to \({\texttt {s}}\)-stage computations. These are computations where the write permission changes \({\texttt {s}}\) times. The problem was first analyzed in [1] and shown to be \({\textsf {NP}} \)-complete for finite state programs. We give matching upper and lower bounds in terms of fine-grained complexity and prove the absence of a polynomial kernel.

Let \({\mathcal {A}}= ( D , a^0, (P_i)_{i \in [1..{\texttt {t}}]})\) be a program. A stage is a computation in \({\mathcal {A}}\) where only one of the threads writes. The remaining threads are restricted to reading the memory. An \({\texttt {s}}\)-stage computation is a computation that can be split into \({\texttt {s}}\) parts, each of which forming a stage. We state the decision problem.


We focus on a parameterization of \({\textsf {BSR}} \) by \({\texttt {P}}\), the maximum number of states of a thread, and \({\texttt {t}}\), the number of threads. Let it be denoted by \({\textsf {BSR}} ({\texttt {P}}, {\texttt {t}})\). We prove that the parameterization is \({\textsf {FPT}} \) and present a matching lower bound. The main result in this section is the absence of a polynomial kernel for \({\textsf {BSR}} ({\texttt {P}}, {\texttt {t}})\). The result is technically involved and shows what makes the problem hard.

Parameterizations of \({\textsf {BSR}} \) involving only \({\texttt {D}}\) and \({\texttt {s}}\) are intractable. We show that \({\textsf {BSR}} \) remains \({\textsf {NP}} \)-hard even if both, \({\texttt {D}}\) and \({\texttt {s}}\), are constants. This proves the existence of an \({\textsf {FPT}} \)-algorithm for those cases unlikely.

Parameterization by Number of States and Threads

We first give an algorithm for \({\textsf {BSR}} \), based on a product construction of automata. Then, we present a lower bound under \({\textsf {ETH}} \). Interestingly, the lower bound shows that we cannot avoid building the product. We conclude with proving the absence of a polynomial kernel. As before, we cross-compose from \({\textsf {3}} \text {-}{\textsf {SAT}} \) but now face the problem that two important parameters in the construction, \({\texttt {P}}\) and \({\texttt {t}}\), are not allowed to depend polynomially on the number of \({\textsf {3}} \text {-}{\textsf {SAT}} \)-instances.

Upper Bound

We show that \({\textsf {BSR}} ({\texttt {P}}, {\texttt {t}})\) is fixed-parameter tractable. The idea is to reduce to reachability on a product automaton. The automaton stores the configurations, the current writer, and counts up to the number of stages \({\texttt {s}}\). To this end, it has \({\mathcal {O}}^*({\texttt {P}}^{{\texttt {t}}})\) many states. Details can be found in the full version of the paper.

Proposition 20

\({\textsf {BSR}} \) can be solved in time \({\mathcal {O}}^*({\texttt {P}}^{\, 2 {\texttt {t}}})\).

Lower Bound

By a reduction from \({\textsf {k}} \times {\textsf {k}} ~ {\textsf {Clique}} \), we show that a \(2^{o({\texttt {t}}\cdot \log ({\texttt {P}}))}\)-time algorithm for \({\textsf {BSR}} \) would contradict \({\textsf {ETH}} \). The above algorithm is optimal.

Proposition 21

\({\textsf {BSR}} \) cannot be solved in time \(2^{o({\texttt {t}}\cdot \log ({\texttt {P}}))}\) unless \({\textsf {ETH}} \) fails.

The reduction constructs from an instance of \({\textsf {k}} \times {\textsf {k}} ~ {\textsf {Clique}} \) an equivalent instance \(({\mathcal {A}}= ( D , a^0, (P_i)_{i\in [1..{\texttt {t}}]}), C^f, {\texttt {s}})\) of \({\textsf {BSR}} \). Moreover, it keeps the parameters small. We have that \({\texttt {P}}= {\mathcal {O}}(k^2)\) and \({\texttt {t}}= {\mathcal {O}}(k)\). As a consequence, a \(2^{o({\texttt {t}}\cdot \log ({\texttt {P}}))}\)-time algorithm for \({\textsf {BSR}} \) would yield an algorithm for \({\textsf {k}} \times {\textsf {k}} ~ {\textsf {Clique}} \) running in time \(2^{o(k \cdot \log (k^2))} = 2^{o(k \cdot \log (k))}\). But this contradicts \({\textsf {ETH}} \).


(Idea) For the reduction, let \(V = [1..k] \times [1..k]\) be the vertices of G. We define \( D = V \cup \lbrace a^0 \rbrace \) to be the domain of the memory. We want the threads to communicate on the vertices of G. For each row we introduce a reader thread \(P_i\) that is responsible for storing a particular vertex of the row. We also add one writer \(P_{ch}\) that is used to steer the communication between the \(P_i\). Our program \({\mathcal {A}}\) is given by the tuple \(( D , a^0, ( (P_i)_{i \in [1..k]}, P_{ch} ))\).

Intuitively, the program proceeds in two phases. In the first phase, each \(P_i\) non-deterministically chooses a vertex from the i-th row and stores it in its state space. This constitutes a clique candidate \((1,j_1), \dots , (k, j_k) \in V\). In the second phase, thread \(P_{ch}\) starts to write a random vertex \((1,j_1')\) of the first row to the memory. The first thread \(P_1\) reads \((1,j_1')\) from the memory and verifies that the read vertex is actually the one from the clique candidate. The computation in \(P_1\) will deadlock if \(j_1' \ne j_1\). The threads \(P_i\) with \(i \ne 1\) also read \((1,j_1')\) from the memory. They have to check whether there is an edge between the stored vertex \((i,j_i)\) and \((1,j_1')\). If this fails in some \(P_i\), the computation in that thread will also deadlock. After this procedure, the writer \(P_{ch}\) guesses a vertex \((2,j_2')\), writes it to the memory, and the verification steps repeat. In the end, after k repetitions of the procedure, we can ensure that the guessed clique candidate is indeed a clique. Formal construction and proof are available in the full version of the paper. \(\square \)

Absence of a Polynomial Kernel

We show that \({\textsf {BSR}} ({\texttt {P}}, {\texttt {t}})\) does not admit a polynomial kernel. To this end, we cross-compose the problem \({\textsf {3}} \text {-}{\textsf {SAT}} \) into \({\textsf {BSR}} ({\texttt {P}}, {\texttt {t}})\).

Theorem 22

\({\textsf {BSR}} ({\texttt {P}}, {\texttt {t}})\) does not admit a polynomial kernel unless \({\textsf {NP}} \subseteq {\textsf {co}} {\textsf {NP}} / {\mathsf {poly}}\).

In the present setting, coming up with a cross-composition is non-trivial. Both parameters, \({\texttt {P}}\) and \({\texttt {t}}\), are not allowed to depend polynomially on the number I of given \({\textsf {3}} \text {-}{\textsf {SAT}} \)-instances. Hence, we cannot construct an NFA that distinguishes the I instances by branching into I different directions. This would cause a polynomial dependence of \({\texttt {P}}\) on I. Furthermore, it is not possible to construct an NFA for each instance as this would cause such a dependence of \({\texttt {t}}\) on I. To circumvent the problems, some deeper understanding of the model is needed.


(Idea) Let \(\varphi _1, \dots , \varphi _I\) be given \({\textsf {3}} \text {-}{\textsf {SAT}} \)-instances, where each two are equivalent under \({{\mathcal {R}}}\), the polynomial equivalence relation of Theorem 10. Then each \(\varphi _\ell \) has m clauses and n variables \(\lbrace x_1, \dots , x_n \rbrace \). We assume \(\varphi _\ell = C^\ell _1 \wedge \dots \wedge C^\ell _m\).

In the program that we construct, the communication is based on 4-tuples of the form \((\ell , j, i, v)\). Intuitively, such a tuple transports the following information: The j-th clause in instance \(\varphi _\ell \), \(C^\ell _j\), can be satisfied by variable \(x_i\) with evaluation v. Hence, our data domain is \( D = ([1..I] \times [1..m] \times [1..n] \times \lbrace 0,1 \rbrace ) \cup \lbrace a^0 \rbrace \).

For choosing and storing an evaluation of the \(x_i\), we introduce so-called variable threads \(P_{x_1}, \dots , P_{x_n}\). In the beginning, each \(P_{x_i}\) non-deterministically chooses an evaluation for \(x_i\) and stores it in its state space.

We further introduce a writer \(P_w\). During a computation, this thread guesses exactly m tuples \((\ell _1, 1, i_1, v_1), \dots , (\ell _m, m, i_m, v_m)\) in order to satisfy m clauses of potentially different instances. Each \((\ell _j, j, i_j, v_j)\) is written to the memory by \(P_w\). All variable threads then start to read the tuple. If \(P_{x_{i}}\) with \(i \ne i_j\) reads it, then the thread will just move one state further since the suggested tuple does not affect the variable \(x_i\). If \(P_{x_i}\) with \(i = i_j\) reads the tuple, the thread will only continue its computation if \(v_j\) coincides with the value that \(P_{x_i}\) guessed for \(x_i\) and, moreover, \(x_i\) with evaluation \(v_j\) satisfies clause \(C^{\ell _j}_{j}\).

Now suppose the writer did exactly m steps while each variable thread did exactly \(m+1\) steps. This proves the satisfiability of m clauses by the chosen evaluation. But these clauses can be part of different instances: It is not ensured that the clauses were chosen from one formula \(\varphi _\ell \). The major difficulty of the cross-composition lies in how to ensure exactly this.

We overcome the difficulty by introducing so-called bit checkers \(P_b\), where \(b \in [1..\log (I)]\). Each \(P_b\) is responsible for the b-th bit of \({\texttt {bin}}(\ell )\), the binary representation of \(\ell \), where \(\varphi _\ell \) is the instance we want to satisfy. When \(P_w\) writes a tuple \((\ell _1, 1, i_1, v_1)\) for the first time, each \(P_b\) reads it and stores either 0 or 1, according to the b-th bit of \({\texttt {bin}}(\ell _1)\). After \(P_w\) has written a second tuple \((\ell _2, 2, i_2, v_2)\), the bit checker \(P_b\) tests whether the b-th bit of \({\texttt {bin}}(\ell _1)\) and \({\texttt {bin}}(\ell _2)\) coincide, otherwise it will deadlock. This will be repeated any time \(P_w\) writes a new tuple to the memory.

Assume the computation does not deadlock in any of the \(P_b\). Then we can ensure that the b-th bit of \({\texttt {bin}}(\ell _j)\) with \(j \in [1..m]\) never changed during the computation. This means that \({\texttt {bin}}(\ell _1) = \dots = {\texttt {bin}}(\ell _m)\). Hence, the writer \(P_w\) has chosen clauses of just one instance \(\varphi _{\ell }\). Moreover, the current evaluation satisfies the formula. Since the parameters are bounded, \({\texttt {P}}\in {\mathcal {O}}(m)\) and \({\texttt {t}}\in {\mathcal {O}}(n + \log (I))\), the construction constitutes a proper cross-composition. For a formal construction and proof, we refer to the full version of the paper. \(\square \)

Variable threads and writer thread are needed for testing satisfiability of clauses. The need for bit checkers comes from ensuring that all clauses stem from the same formula. We illustrate the notion with an example.

Example 23

Let four formulas \(\varphi _1, \varphi _2, \varphi _3, \varphi _4\) with two clauses each be given. We show how the bit checkers are constructed. To this end, we first encode the index of the instances as binary numbers using two bits. The encoding is shown in Fig. 4 on the right hand side. Note the offset by one in the encoding.

We focus on the bit checker \(P_{b_1}\) responsible for the first bit. It is illustrated in Fig. 4 on the left hand side. Note that the label \(\ell = 1, \ell = 3\) refers to transitions of the form \(?(\ell ,j,i,v)\) with \(\ell \) either 1 or 3 and arbitrary values for i,j, and v. On reading the first of these tuples, \(P_{b_1}\) stores the first bit of \(\ell \) in its state space. The blue marked states store that \(b_1 = 0\), the red states store \(b_1 = 1\). Then, the bit checker can only continue on reading tuples \((\ell ,j,i,v)\) where the first bit of \(\ell \) matches the stored bit. In the case of \(b_1 = 0\), this means that \(P_{b_1}\) can only read tuples \((\ell ,j,i,v)\) with \(\ell \) either 1 or 3.

Assume the writer thread has output two tuples \((\ell _1,1,i_1,v_1)\) and \((\ell _2,2,i_2,v_2)\) and the bit checker \(P_{b_1}\) has reached a last state. Since the computation did not deadlock on \(P_{b_1}\), we know that the first bits of \(\ell _1\) and \(\ell _2\) coincide. If the bit checker for the second bit does not deadlock as well, we get that \(\ell _1 = \ell _2\). Hence, the writer has chosen two clauses from one instance \(\varphi _{\ell _1}\).

Fig. 4

A binary encoding (right) of the numbers 1 up to 4 using two bits. First bits are either marked blue, if they are 0, or red if they are 1. The bit checker \(P_{b_1}\) (left) focuses on the first bit. The label \(\ell = 1, \ell = 3\) means that \(P_{b_1}\) has transitions on \(?(\ell ,j,i,v)\) for \(\ell = 1,3\) and arbitrary values for ij, and v. The blue marked states store that the first bit \(b_1\) is 0. Red marked states store that \(b_1\) is 1. (Color figure online)


We show that parameterizations of \({\textsf {BSR}} \) involving only \({\texttt {s}}\) and \({\texttt {D}}\) are intractable. To this end, we prove that \({\textsf {BSR}} \) remains \({\textsf {NP}} \)-hard even if both parameters are constant. This is surprising as the number of stages \({\texttt {s}}\) seems to be a powerful parameter. Introducing such a bound in simultaneous reachability lets the complexity drop from \({\textsf {PSPACE}} \) to \({\textsf {NP}} \). But still, it is not enough to guarantee an \({\textsf {FPT}} \)-algorithm.

Proposition 24

\({\textsf {BSR}} \) is \({\textsf {NP}} \)-hard even if both \({\texttt {s}}\) and \({\texttt {D}}\) are constant.

The proposition implies intractability: Assume there is an \({\textsf {FPT}} \)-algorithm A for \({\textsf {BSR}} \) running in time \(f({\texttt {s}},{\texttt {D}}) \cdot poly(|x|)\), where x denotes the input. Then \({\textsf {BSR}} '\), the variant of \({\textsf {BSR}} \) where \({\texttt {s}}\) and \({\texttt {D}}\) are constant, can also be solved by A. In this case, the runtime of A is \({\mathcal {O}}(poly(|x|))\) since \(f({\texttt {s}},{\texttt {D}})\) is a constant on every instance of \({\textsf {BSR}} '\). But this contradicts the \({\textsf {NP}} \)-hardness of \({\textsf {BSR}} '\) which is shown in the proposition.


(Idea) We give a reduction from \({\textsf {3}} \text {-}{\textsf {SAT}} \) to \({\textsf {BSR}} \) that keeps both parameters constant. Let \(\varphi \) be a \({\textsf {3}} \text {-}{\textsf {SAT}} \)-instance with m clauses and variables \(x_1, \dots , x_n\). We construct a program \({\mathcal {A}}= ( D , a^0, P_1, \dots , P_n, P_v)\) with \({\texttt {D}}= 4\) different memory symbols that can only run 1-stage computations.

The program cannot communicate on literals directly, as this would cause a blow-up in parameter \({\texttt {D}}\). Instead, variables and evaluations are encoded in binary in the following way. Let \(\ell \) be a literal in \(\varphi \). It consists of a variable \(x_i\) and an evaluation \(v \in \lbrace 0,1 \rbrace \). The padded binary encoding \({\texttt {bin}}_\#(i) \in (\lbrace 0,1 \rbrace . \#)^{\log (n)+1}\) of i is the usual binary encoding where each bit is separated by a \(\#\). The string \({\mathtt{Enc}}(\ell ) = v \# {\texttt {bin}}_\#(i)\) encodes that variable \(x_i\) has evaluation v. We need the padding symbol \(\#\) to prevent the threads in \({\mathcal {A}}\) from reading the same symbol more than once. Program \({\mathcal {A}}\) communicates by passing messages of the form \({\mathtt{Enc}}(\ell )\). To this end, we need the data domain \( D = \lbrace a^0,\#,0,1 \rbrace \).

The program contains threads \(P_i, i \in [1..n]\), called variable threads. Initially, these threads choose an evaluation for the variables and store it: Each \(P_i\) can branch on reading \(a^0\) and choose whether it assigns 0 or 1 to \(x_i\). Then, a verifier thread \(P_v\) starts to iterate over the clauses. For each clause C, it picks a literal \(\ell \in C\) that should evaluate to true and writes its encoding \({\mathtt{Enc}}(\ell )\) to the memory. Each of the \(P_i\) reads \({\mathtt{Enc}}(\ell )\). Note that reading and writing \({\mathtt{Enc}}(\ell )\) needs a sequence of transitions. In the construction, we ensure that all the needed states and transitions are provided. It is the task of each \(P_i\) to check whether the chosen literal \(\ell \) is conform with the chosen evaluation for \(x_i\). To this end, we distinguish two cases.

  1. (1)

    If \(\ell \) involves a variable \(x_j\) with \(j \ne i\), variable thread \(P_i\) just continues its computation by reading the whole string \({\mathtt{Enc}}(\ell )\).

  2. (2)

    If \(\ell \) involves \(x_i\), \(P_i\) has to ensure that the stored evaluation coincides with the one sent by the verifier. To this end, \(P_i\) can only continue its computation if the first bit in \({\mathtt{Enc}}(\ell )\) shows the correct evaluation. Formally, there is only an outgoing path of transitions on \({\mathtt{Enc}}(x_i)\) if \(P_i\) stored 1 as evaluation and on \({\mathtt{Enc}}(\lnot x_i)\) if it stored 0.

Note that each time \(P_v\) picks a literal \(\ell \), all \(P_i\) read \({\mathtt{Enc}}(\ell )\), even if the literal involves a different variable. This means that the \(P_i\) count how many literals have been seen already. This is important for correctness: The threads will only terminate if they have read a word of fixed length and did not miss a single symbol. Phrased differently, there is no loss in the communication between \(P_v\) and the \(P_i\).

Now assume \(P_v\) iterated through all m clauses and none of the variable threads got stuck. Then, each of them read exactly m encodings without running into a deadlock. Hence, the picked literals were all conform with the evaluation chosen by the \(P_i\). This means that a satisfying assignment for \(\varphi \) is found.

During a computation of \({\mathcal {A}}\), the verifier \(P_v\) is the only thread that has write permission. Hence, each computation of \({\mathcal {A}}\) consists of a single stage. For a formal construction, we refer to the full version of the paper. \(\square \)


We have studied several parameterizations of \({\textsf {LCR}} \) and \({\textsf {BSR}} \), two safety verification problems for shared memory concurrent programs. In \({\textsf {LCR}} \), a designated leader thread interacts with a number of equal contributor threads. The task is to decide whether the leader can reach an unsafe state. The problem \({\textsf {BSR}} \) is a generalization of bounded context switching. A computation gets split into stages, periods where writing is restricted to one thread. Then, \({\textsf {BSR}} \) asks whether all threads can reach a final state simultaneously during an \({\texttt {s}}\)-stage computation.

For \({\textsf {LCR}} \), we identified the size of the data domain \({\texttt {D}}\), the size of the leader \({\texttt {L}}\) and the size of the contributors \({\texttt {C}}\) as parameters. Our first algorithm showed that \({\textsf {LCR}} ({\texttt {D}}, {\texttt {L}})\) is \({\textsf {FPT}} \). Then we modified the algorithm to obtain a verification procedure valuable for practical instances. The main insight was that due to a factorization along strongly connected components, the impact of \({\texttt {L}}\) can be reduced to a polynomial factor in the time complexity. We also proved the absence of a polynomial kernel for \({\textsf {LCR}} ({\texttt {D}}, {\texttt {L}})\) and presented an \({\textsf {ETH}} \)-based lower bound which shows that the upper bound is a root-factor away from being optimal.

For \({\textsf {LCR}} ({\texttt {C}})\) we presented a dynamic programming, running in \({\mathcal {O}}^*(2^{\texttt {C}})\) time. The algorithm is based on slice-wise reachability. This reduces a reachability problem on a large graph to reachability problems on subgraphs (slices) that are solvable in polynomial time. Moreover, we gave a tight lower bound based on \({\textsf {Set~Cover}} \) and proved the absence of a polynomial kernel.

Parameterizations different from \({\textsf {LCR}} ({\texttt {D}},{\texttt {L}})\) and \({\textsf {LCR}} ({\texttt {C}})\) were shown to be intractable. We gave reductions from \({\textsf {k \text {-} Clique}} \) and proved \({\textsf {W}} [1]\)-hardness.

The parameters of interest for \({\textsf {BSR}} \) are the maximum size of a thread \({\texttt {P}}\) and the number of threads \({\texttt {t}}\). We have shown that a parameterization by both parameters is \({\textsf {FPT}} \) and gave a matching lower bound. The main contribution was to prove it unlikely that a polynomial kernel exists for \({\textsf {BSR}} ({\texttt {P}}, {\texttt {t}})\). The proof relies on a technically involved cross-composition that avoids a polynomial dependence of the parameters on the number of given \({\textsf {3}} \text {-}{\textsf {SAT}} \)-instances.

Parameterizations involving other parameters like \({\texttt {s}}\) or \({\texttt {D}}\) were proven to be intractable for \({\textsf {BSR}} \). We gave an \({\textsf {NP}} \)-hardness proof where \({\texttt {s}}\) and \({\texttt {D}}\) are constant.

Extension of the Model

In this work, the considered model for programs allows the memory to consist of a single cell. We discuss whether the presented results carry over when the number of memory cells increases. Having multiple memory cells is referred to as supporting global variables. Extending the definition of programs in Sect. 2 to global variables is straightforward.

For the problem \({\textsf {LCR}} \), allowing global variables is a rather powerful mechanism. Let \({\textsf {LCR}} _ Var \) denote the problem \({\textsf {LCR}} \) where the input is a program featuring global variables. The interesting parameters for the problem are \({\texttt {D}}\), \({\texttt {L}}\), \({\texttt {C}}\), and \({\texttt {v}}\), the number of global variables. It turns out that \({\textsf {LCR}} _ Var \) is \({\textsf {PSPACE}} \)-hard, even when \({\texttt {C}}\) is constant. One can reduce the intersection emptiness problem for finite automata to \({\textsf {LCR}} _ Var \). The reduction makes use only of the leader, contributors are not needed.

A program \({\mathcal {A}}\) with global variables can always be reduced to a program \({\mathcal {A}}'\) with a single memory cell [26]. Roughly, the reduction constructs the leader of \({\mathcal {A}}'\) in such a way that it can store the memory contents of \({\mathcal {A}}\) and manage contributor accesses to the memory. This means the new leader needs exponentially many states since there are \({\texttt {D}}^{\texttt {v}}\) many possible memory valuations. The domain and the contributor of \({\mathcal {A}}'\) are of polynomial size. In fact, we can then apply the algorithm from Sect. 3.1 to the program \({\mathcal {A}}'\). The runtime depends exponentially only on the parameters \({\texttt {D}}\), \({\texttt {L}}\), and \({\texttt {v}}\). This shows that \({\textsf {LCR}} _ Var ({\texttt {D}}, {\texttt {L}}, {\texttt {v}})\) is fixed-parameter tractable. It is an interesting question whether this algorithm can be improved. Moreover, it is open whether there are other parameterizations of \({\textsf {LCR}} _ Var \) that have an \({\textsf {FPT}} \)-algorithm. A closer investigation is considered future work.

For \({\textsf {BSR}} \), allowing global variables also leads to \({\textsf {PSPACE}} \)-hardness. The problem \({\textsf {BSR}} _ Var \), defined similarly to \({\textsf {LCR}} _ Var \), is \({\textsf {PSPACE}} \)-hard already for a constant number of threads. In fact, the proof is similar to the hardness of \({\textsf {LCR}} _ Var \) where only one thread is needed. To obtain an algorithm for the problem, we modify the construction from Proposition 20. The resulting product automaton then also maintains the values of the global variables. This shows membership in \({\textsf {PSPACE}} \). But the size of the product now also depends exponentially on \({\texttt {D}}\) and \({\texttt {v}}\). The interesting question is whether we can find an algorithm that avoids an exponential dependence on one of the parameters \({\texttt {P}}, {\texttt {t}}, {\texttt {D}}\) or \({\texttt {v}}\). It is a matter of future work to examine the precise complexity of the different parameterizations.


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Chini, P., Meyer, R. & Saivasan, P. Fine-Grained Complexity of Safety Verification. J Autom Reasoning 64, 1419–1444 (2020).

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  • Parameterized verification
  • Parameterized complexity
  • Fine-grained complexity
  • Safety verification