Reachability Analysis for Termination and Confluence of Rewriting
Abstract
In term rewriting, reachability analysis is concerned with the problem of deciding whether or not one term is reachable from another by rewriting. Reachability analysis has several applications in termination and confluence analysis of rewrite systems. We give a unified view on reachability analysis for rewriting with and without conditions by means of what we call reachability constraints. Moreover, we provide several techniques that fit into this general framework and can be efficiently implemented. Our experiments show that these techniques increase the power of existing termination and confluence tools.
Keywords
Reachability analysis Termination Confluence Conditional term rewriting Infeasibility1 Introduction
Reachability analysis for term rewriting [6] is concerned with the problem of, given a rewrite system Open image in new window , a source term s and a target term t, deciding whether the source reduces to the target by rewriting, which is usually written Open image in new window . A useful generalization of this problem is the (un)satisfiability of the following reachability problem: given terms s and t containing variables, decide whether there is a substitution \(\sigma \) such that Open image in new window or not. This problem, also called (in)feasibility by Lucas and Guitiérrez [11], has various applications in termination and confluence analysis for plain and conditional rewriting.
This can be understood as a form of safety analysis, as illustrated below.
Example 1
In termination analysis we are typically interested in unsatisfiability of reachability and can thereby rule out certain recursive calls as potential source of nontermination. For confluence analysis of conditional term rewriting, infeasibility is crucial: some other techniques do not apply before critical pairs are shown infeasible, and removal of infeasible rules simplifies proofs.
In this work we provide a formal framework that allows us to uniformly speak about (un)satisfiability of reachability for plain and conditional rewriting, and give several techniques that are useful in practice.

We introduce the syntax and semantics of reachability constraints (Sect. 3) and formulate their satisfiability problem. We recast several concrete techniques for reachability analysis in the resulting framework.

We present a new, simple, and efficient technique for reachability analysis based on what we call the symbol transition graph of a rewrite system (Sect. 4.1) and extend it to conditional rewriting (Sect. 5.2).

Additionally, we generalize the prevalent existing technique for term rewriting to what we call lookahead reachability (Sect. 4.2) and extend it to the conditional case (Sect. 5.3).

Then, we present a new result for conditional rewriting that is useful for proving conditional rules infeasible (Sect. 5.1).

Finally, we evaluate the impact of our work on existing automated tools NaTT [16] and ConCon [13] (Sect. 6).
2 Preliminaries
In the remainder, we assume some familiarity with term rewriting. Nevertheless, we recall required concepts and notations below. For further details on term rewriting, we refer to standard textbooks [3, 14].
Throughout the paper Open image in new window denotes a set of function symbols with associated arities, and \(\mathcal {V}\) a countably infinite set of variables (so that fresh variables can always be picked) such that Open image in new window . A term is either a variable Open image in new window or of the form \(f(t_1,\dots ,t_n)\), where n is the arity of Open image in new window and the arguments \(t_1,\dots ,t_n\) are terms. The set of all terms over Open image in new window and \(\mathcal {V}\) is denoted by Open image in new window . The set of variables occurring in a term t is denoted by \(\mathsf {Var}(t)\). The root symbol of a term \(t = f(t_1,\dots ,t_n)\) is f and denoted by \(\mathsf {root}(t)\). When we want to indicate that a term is not a variable, we sometimes write \(f(...)\), where “\(...\)” denotes an arbitrary list of terms.
A substitution is a mapping Open image in new window . Given a term t, \(t\sigma \) denotes the term obtained by replacing every occurrence of variable x in t by \(\sigma (x)\). The domain of a substitution \(\sigma \) is Open image in new window , and \(\sigma \) is idempotent if Open image in new window for every Open image in new window . A renaming is a bijection Open image in new window . Two terms s and t are unifiable if \(s\sigma = t\sigma \) for some substitution \(\sigma \), which is called a unifier of s and t.
A context is a term with exactly one occurrence of the special symbol Open image in new window . We write C[t] for the term resulting from replacing Open image in new window in context C by term t.
A rewrite rule is a pair of terms, written Open image in new window , such that the variable conditions Open image in new window and \(\mathsf {Var}(l) \supseteq \mathsf {Var}(r)\) hold. By a variant of a rewrite rule we mean a rule that is obtained by consistently renaming variables in the original rule to fresh ones. A term rewrite system (TRS) is a set Open image in new window of rewrite rules. A function symbol Open image in new window is defined in Open image in new window if Open image in new window , and the set of defined symbols in Open image in new window is Open image in new window . We call Open image in new window a constructor.
There is an Open image in new window rewrite step from s to t, written Open image in new window , iff there exist a context C, a substitution \(\sigma \), and a rule Open image in new window such that \(s = C[l\sigma ]\) and \(t = C[r\sigma ]\). We write Open image in new window if Open image in new window (called a root step), and Open image in new window (called a nonroot step), otherwise. We say a term \(s_0\) is Open image in new window terminating if it starts no infinite rewrite sequence Open image in new window , and say Open image in new window is terminating if every term is Open image in new window terminating.
For a relation \({\rightarrowtail } \subseteq A \times A\), we denote its transitive closure by \(\rightarrowtail ^+\) and reflexive transitive closure by \(\rightarrowtail ^*\). We say that Open image in new window are joinable (meetable) at Open image in new window with respect to \(\rightarrowtail \) if \(a_i \rightarrowtail ^* b\) (\(b \rightarrowtail ^* a_i\)) for every Open image in new window .
3 Reachability Constraint Satisfaction
In this section we introduce the syntax and semantics of reachability constraints, a framework that allows us to unify several concrete techniques for reachability analysis on an abstract level. Reachability constraints are firstorder formulas^{1} with a single binary predicate symbol whose intended interpretation is reachability by rewriting with respect to a given rewrite system.
Definition 1
To save some space, we use conventional notation like Open image in new window and Open image in new window .
As mentioned above, the semantics of reachability constraints is defined with respect to a given rewrite system. In the following we define satisfiability of constraints with respect to a TRS. (This definition will be extended to conditional rewrite systems in Sect. 5).
Definition 2

Open image in new window if Open image in new window or Open image in new window ;

Open image in new window if Open image in new window and Open image in new window ;

Open image in new window if Open image in new window does not hold;

Open image in new window if Open image in new window for every \(\sigma '\) that coincides with \(\sigma \) on Open image in new window .

Open image in new window if Open image in new window for some \(\sigma '\) that coincides with \(\sigma \) on Open image in new window .
We say \(\phi \) and \(\psi \) are equivalent modulo Open image in new window , written Open image in new window , when Open image in new window iff Open image in new window for all \(\sigma \). We say \(\phi \) and \(\psi \) are (logically) equivalent, written \(\phi \,\equiv _{}\,\psi \), if they are equivalent modulo any Open image in new window . We say \(\phi \) is satisfiable modulo Open image in new window , written Open image in new window , if there is a substitution \(\sigma \) that satisfies \(\phi \) modulo Open image in new window , and call \(\sigma \) a solution of \(\phi \) with respect to Open image in new window .
Checking for satisfiability of reachability constraints is for example useful for proving termination of term rewrite systems via the dependency pair method [2], or more specifically in dependency graph analysis. For the dependency pair method, we assume a fresh marked symbol \(f^\sharp \) for every Open image in new window , and write \(s^\sharp \) to denote the term \(f^\sharp (s_1,\dots ,s_n)\) for \(s = f(s_1,\dots ,s_n)\). The set of dependency pairs of a TRS Open image in new window is Open image in new window . The standard definition of the dependency graph of a TRS [2] can be recast using reachability constraints as follows:
Definition 3
(Dependency Graph). Given a TRS Open image in new window , its dependency graph Open image in new window is the directed graph over Open image in new window where there is an edge from Open image in new window to Open image in new window iff Open image in new window , where \(\alpha \) is a renaming of variables such that Open image in new window .
The nodes of the dependency graph correspond to the possible recursive calls in a program (represented by a TRS), while its edges encode the information which recursive calls can directly follow each other in arbitrary program executions. This is the reason why dependency graphs are useful for investigating the termination behavior of TRSs, as captured by the following result.
Theorem 1
([10]). A TRS Open image in new window is terminating iff for every strongly connected component \(\mathcal {C}\) of an over approximation of Open image in new window , there is no infinite chain Open image in new window where every \(t_i\) is Open image in new window terminating.
Example 2
The most popular method today for checking reachability during dependency graph analysis is unifiability between the target and an approximation of the topmost part of the source (its “cap”) that does not change under rewriting, which is computed by the \(\mathsf {tcap}_{\mathcal {R}}\) function [9].
Definition 4
\(\mathbf{(}{\mathsf {tcap}_{}}\mathbf{).}\) Let Open image in new window be a TRS. We recursively define \(\mathsf {tcap}_{\mathcal {R}}(t)\) for a given term t as follows: \(\mathsf {tcap}_{\mathcal {R}}(x)\) is a fresh variable if Open image in new window ; \(\mathsf {tcap}_{\mathcal {R}}(f(t_1,\dots ,t_n))\) is a fresh variable if \(u = f(\mathsf {tcap}_{\mathcal {R}}(t_1),\dots ,\mathsf {tcap}_{\mathcal {R}}(t_n))\) unifies with some lefthand side of the rules in Open image in new window ; otherwise, it is u.
The standard way of checking for nonreachability that is implemented in most tools is captured by of the following proposition.
Proposition 1
If \(\mathsf {tcap}_{\mathcal {R}}(s)\) and t are not unifiable, then Open image in new window .
Example 3
Proposition 1 cannot prove the unsatisfiability of (2) of Example 2, since the term cap of the source Open image in new window , where Open image in new window , Open image in new window , Open image in new window are fresh variables, is unifiable with the target Open image in new window .
4 Reachability in Term Rewriting
In this section we introduce some techniques for analyzing (un)satisfiability of reachability constraints. The first one described below formulates an obvious observation: no root rewrite step is applicable when starting from a term whose root is a constructor.
Definition 5

Open image in new window if \(f \ne g\), and
The intention of nonroot reachability constraints is to encode zero or more steps of nonroot rewriting, in the following sense.
Lemma 1
For Open image in new window , Open image in new window iff Open image in new window .
Proof
The claim vacuously follows if \(\mathsf {root}(s) \ne \mathsf {root}(t)\). So let \(s = f(s_1,\dots ,s_n)\) and \(t = f(t_1,\dots ,t_n)\). We have Open image in new window iff Open image in new window iff Open image in new window . \(\square \)
Combined with the observation that no root step is applicable to a term whose root symbol is a constructor, we obtain the following reformulation of a folklore result that reduces reachability to direct subterms.
Proposition 2
If \(s = f(...)\) with Open image in new window and \(t \notin \mathcal {V}\), then Open image in new window .
Proposition 2 is directly applicable in the analysis of dependency graphs.
Example 4
4.1 Symbol Transition Graphs
Here we introduce a new, simple and efficient way of overapproximating reachability by tracking the relation of root symbols of terms according to a given set of rewrite rules. We first illustrate the intuition by an example.
Example 5

If Open image in new window since nonroot steps preserve the root symbol and no root steps are applicable to terms of the form Open image in new window .

If Open image in new window , then Open image in new window since nonroot steps preserve the root symbol and the only possible root step is Open image in new window .

If Open image in new window , then Open image in new window by the same reasoning.

If Open image in new window , then t can be any term and \(\mathsf {root}(t)\) can be arbitrary.
This informal argument is captured by the following definition.
Definition 6
(Symbol Transition Graphs). The symbol transition graph Open image in new window of a TRS Open image in new window over signature Open image in new window is the graph Open image in new window , where \(f \rightarrowtail _{\mathcal {R}}g\) iff Open image in new window contains a rule of form Open image in new window or Open image in new window with Open image in new window .
The following result tells us that for nonvariable terms the symbol transition graph captures the relation between the root symbols of root rewrite steps.
Lemma 2
Proof
By assumption there exist Open image in new window and \(\sigma \) such that \(s = l\sigma \) and \(r\sigma = t\). If Open image in new window then either Open image in new window or \(\mathsf {root}(s) = \mathsf {root}(l) \rightarrowtail _{\mathcal {R}}\mathsf {root}(t)\). Otherwise, \(\mathsf {root}(s) = \mathsf {root}(l) \rightarrowtail _{\mathcal {R}}\mathsf {root}(r) = \mathsf {root}(t)\). \(\square \)
Since every rewrite sequence is composed of subsequences that take place entirely below the root (and hence do not change the root symbol) separated by root steps, we can extend the previous result to rewrite sequences.
Lemma 3
If Open image in new window then \(f \rightarrowtail _{\mathcal {R}}^* g\).
Proof
We prove the claim for arbitrary s and f by induction on the derivation length of Open image in new window . The base case is trivial, so consider Open image in new window . Since Open image in new window , we have Open image in new window with \(s' = f'(...)\). Thus the induction hypothesis yields \(f' \rightarrowtail _{\mathcal {R}}^* g\). If Open image in new window then by Lemma 2 we conclude \(f \rightarrowtail _{\mathcal {R}}f' \rightarrowtail _{\mathcal {R}}^* g\), and otherwise \(f = f' \rightarrowtail _{\mathcal {R}}^* g\). \(\square \)
It is now straightforward to derive the following from Lemma 3.
Corollary 1
If \(f \rightarrowtail _{\mathcal {R}}^* g\) does not hold, then Open image in new window .
Example 6
The symbol transition graph for Example 5 is depicted in Fig. 1(a). By Corollary 1 we can conclude, for instance, Open image in new window is unsatisfiable.
Corollary 1 is useful for checking (un)satisfiability of Open image in new window , only if neither s nor t is a variable. However, the symbol transition graph is also useful for unsatisfiability in the case when s and t may be variables.
Proposition 3
If Open image in new window for \(t_1 = g_1(...), \,\ldots , t_n = g_n(...)\), then \(g_1, \ldots , g_n\) are meetable with respect to \(\rightarrowtail _{\mathcal {R}}\).
Proof
By assumption there is a substitution \(\sigma \) such that Open image in new window . Clearly Open image in new window is not possible. Thus, suppose \(x\sigma = f(...)\) for some f. Finally, from Lemma 3, we have \(f \rightarrowtail _{\mathcal {R}}^* g_1, \ldots , f \rightarrowtail _{\mathcal {R}}^* g_n\) and thereby conclude that \(g_1, \ldots , g_n\) are meetable at f. \(\square \)
The dual of Proposition 3 is proved in a similar way, but with some special care to ensure Open image in new window .
Proposition 4
If Open image in new window for \(s_1 = f_1(...), \ldots , s_n = f_n(...)\), then \(f_1, \ldots , f_n\) are joinable with respect to \(\rightarrowtail _{\mathcal {R}}\).
Example 7
(Continuation of Example 4). Due to Proposition 3, proving (3) unsatisfiable reduces to proving that Open image in new window are not meetable with respect to \(\rightarrowtail _{\mathcal {R}}\). This is obvious from the symbol transition graph depicted in Fig. 1(b). Hence, we conclude the termination of Open image in new window .
Example 8
4.2 LookAhead Reachability
Here we propose another method for overapproximating reachability, which eventually subsumes the \(\mathsf {tcap}_{}\)unifiability method when target terms are linear. Note that this condition is satisfied in the dependency graph approximation of leftlinear TRSs. Our method is based on the observation that any rewrite sequence either contains at least one root step, or takes place entirely below the root. This observation can be captured using our reachability constraints.
Definition 7
In the definition above, the intuition is that if there are any root steps inside a rewrite sequence then we can pick the first one, which is only preceded by nonroot steps. The following theorem justifies this intuition.
Theorem 2
If Open image in new window , then Open image in new window .
Proof

Open image in new window , that is, \(l = f(l_1,\dots ,l_n)\) and Open image in new window ;

Open image in new window , that is, Open image in new window .
In combination, we have Open image in new window .

No root step is involved: Open image in new window . Then Lemma 1 implies Open image in new window .

At least one root step is involved: there is a rule Open image in new window and a substitution \(\theta \) such that Open image in new window and Open image in new window . Since variables in \(l\theta \) must occur in \(s\sigma \) (due to our assumptions on rewrite rules), we have \(l\theta = l\theta \sigma \) since \(\sigma \) is idempotent. Thus from Lemma 1 we have Open image in new window . Further, variables in \(r\theta \) must occur in \(l\theta \) and thus in \(s\theta \), we also have Open image in new window , and hence Open image in new window . This concludes Open image in new window . \(\square \)
Proposition 2 is a corollary of Theorem 2 together with the following easy lemma, stating that if the root symbol of the source term is not a defined symbol, then no root step can occur.
Lemma 4
Example 9
Disjuncts 1, 3, and 4 expand to Open image in new window by definition of Open image in new window . For disjunct 2, applying Theorem 2 or Proposition 2 to Open image in new window yields Open image in new window .
Note that Theorem 2 can be applied arbitrarily often. Thus, to avoid nontermination in an implementation, we need to control how often it is applied. For this purpose we introduce the following definition.
Definition 8
It easily follows from Theorem 2 and induction on k that the kfold lookahead preserves the semantics of reachability constraints.
Corollary 2
The following results indicate that, whenever \(\mathsf {tcap}_{\mathcal {R}}\)unifiability (Proposition 1) proves Open image in new window unsatisfiable for linear t, \(\mathsf {L}_{\mathcal {R}}^{1}\) can also conclude it.
Lemma 5
Let \(s = f(s_1,\dots ,s_n)\) and Open image in new window be a linear term, and suppose that \(f(\mathsf {tcap}_{\mathcal {R}}(s_1),\dots ,\mathsf {tcap}_{\mathcal {R}}(s_n))\) does not unify with t or any lefthand side in Open image in new window . Then Open image in new window .
Proof
By structural induction on s. First, we show Open image in new window . This is trivial if \(\mathsf {root}(t) \ne f\). So let \(t = f(t_1,\dots ,t_n)\). By assumption there is an Open image in new window such that \(\mathsf {tcap}_{\mathcal {R}}(s_i)\) does not unify with \(t_i\). Hence \(\mathsf {tcap}_{\mathcal {R}}(s_i)\) cannot be a fresh variable, and thus \(s_i\) is of the form \(g(u_1,\dots ,u_m)\) and \(\mathsf {tcap}_{\mathcal {R}}(s_i) = g(\mathsf {tcap}_{\mathcal {R}}(u_1),\dots ,\mathsf {tcap}_{\mathcal {R}}(u_m))\) is not unifiable with any lefthand side in Open image in new window . Therefore, the induction hypothesis applies to \(s_i\), yielding Open image in new window . This concludes Open image in new window .
Second, we show Open image in new window . To this end, we show for an arbitrary variant Open image in new window of a rule in Open image in new window that Open image in new window . This is clear if \(\mathsf {root}(l) \ne f\). So let \(l = f(l_1,\dots ,l_n)\). By assumption there is an Open image in new window such that \(\mathsf {tcap}_{\mathcal {R}}(s_i)\) and \(l_i\) are not unifiable. By a similar reasoning as above the induction hypothesis applies to \(s_i\) and yields Open image in new window . This concludes Open image in new window . \(\square \)
Corollary 3
If \(\mathsf {tcap}_{\mathcal {R}}(s)\) and t are not unifiable, then Open image in new window .
5 Conditional Rewriting
Conditional rewriting is a flavor of rewriting where rules are guarded by conditions. On the one hand, this gives us a boost in expressiveness in the sense that it is often possible to directly express equations with preconditions and that it is easier to directly express programming constructs like the whereclauses of Haskell. On the other hand, the analysis of conditional rewrite systems is typically more involved than for plain rewriting.
In this section we first recall the basics of conditional term rewriting. Then, we motivate the importance of reachability analysis for the conditional case. Finally, we extend the techniques of Sect. 4 to conditional rewrite systems.
Preliminaries. A conditional rewrite rule Open image in new window consists of two terms Open image in new window and r (the lefthand side and righthand side, respectively) and a list \(\phi \) of pairs of terms (its conditions). A conditional term rewrite system (CTRS for short) is a set of conditional rewrite rules. Depending on the interpretation of conditions, conditional rewriting can be separated into several classes. For the purposes of this paper we are interested in oriented CTRSs, where conditions are interpreted as reachability constraints with respect to conditional rewriting. Hence, from now on we identify conditions Open image in new window with the reachability constraint Open image in new window , and the empty list with Open image in new window (omitting “ Open image in new window ” from rules).
Definition 9
(Level Satisfiability). Let Open image in new window be a CTRS and \(\phi \) a reachability constraint. We say that a substitution \(\sigma \) satisfies \(\phi \) modulo Open image in new window at level i, whenever Open image in new window . If we are not interested in a specific satisfying substitution we say that \(\phi \) is satisfiable modulo Open image in new window at level i and write Open image in new window (or just \(\mathsf {SAT}_{i}(\phi )\) whenever Open image in new window is clear from the context).
5.1 Infeasibility
The main area of interest for reachability analysis in the conditional case is checking for infeasibility. While a formal definition of this concept follows below, for the moment, think of it as unsatisfiability of conditions. The two predominant applications of infeasibility are: (1) if the conditions of a rule are unsatisfiable, the rule can never be applied and thus safely be removed without changing the induced rewrite relation; (2) if the conditions of a conditional critical pair (which arises from confluence analysis of CTRSs) are unsatisfiable, then it poses no problem to confluence and can safely be ignored.
Definition 10
(Infeasibility). We say that a conditional rewrite rule Open image in new window is applicable at level i with respect to a CTRS Open image in new window iff Open image in new window . A set \(\mathcal {S}\) of rules is infeasible with respect to Open image in new window when no rule in \(\mathcal {S}\) is applicable at any level.
The next theorem allows us to remove some rules from a CTRS while checking for infeasibility of rules.
Theorem 3
A set \(\mathcal {S}\) of rules is infeasible with respect to a CTRS Open image in new window iff it is infeasible with respect to Open image in new window .
Proof
The ‘only if’ direction is trivial. Thus we concentrate on the ‘if’ direction. To this end, assume that \(\mathcal {S}\) is infeasible with respect to Open image in new window , but not infeasible with respect to Open image in new window . That is, at least one rule in \(\mathcal {S}\) is applicable at some level with respect to Open image in new window . Let m be the minimum level such that there is a rule Open image in new window that is applicable at level m with respect to Open image in new window . Now if \(m = 0\) then Open image in new window is applicable at level 0 and thus Open image in new window , which trivially implies Open image in new window , contradicting the assumption that all rules in \(\mathcal {S}\) are infeasible with respect to Open image in new window . Otherwise, \(m = k+1\) for some \(k \ge 0\) and since Open image in new window is applicable at level m we have Open image in new window . Moreover, the rewrite relations Open image in new window and Open image in new window coincide (since all rules in \(\mathcal {S}\) are infeasible at levels smaller than m by our choice of m). Thus we also have Open image in new window , again contradicting the assumption that all rules in \(\mathcal {S}\) are infeasible with respect to Open image in new window . \(\square \)
The following example from the confluence problems data base (Cops)^{3} shows that Theorem 3 is beneficial for showing infeasibility of conditional rewrite rules.
Example 10
5.2 Symbol Transition Graphs in the Presence of Conditions
In the presence of conditions in rules we replace Definition 6 by the following inductive definition:
Definition 11
The example below shows the difference between the symbol transition graph for TRSs (which can be applied as a crude overapproximation also to CTRSs by dropping all conditions) and the inductive symbol transition graph for CTRSs.
Example 11
The inductive symbol transition graph gives us a sufficient criterion for concluding nonreachability with respect to a given CTRS, as shown in the following.
Lemma 6
If Open image in new window then \(f \rightarrowtail _{\mathcal {R}}^* g\).
Proof
Let \(s = f(...)\) and \(u = g(...)\) and assume that s rewrites to u at level i, that is, Open image in new window . We prove the statement by induction on the level i. If \(i = 0\) then we are done, since Open image in new window is empty and therefore \(f(...) = s = u = g(...)\), which trivially implies \(f \rightarrowtail _{\mathcal {R}}^* g\). Otherwise, \(i = j + 1\) and we obtain the induction hypothesis (IH) that Open image in new window implies \(\mathsf {root}(s) \rightarrowtail _{\mathcal {R}}^* \mathsf {root}(t)\) for arbitrary nonvariable terms s and t. We proceed to show that Open image in new window implies \(f \rightarrowtail _{\mathcal {R}}^* g\) by an inner induction on the length of this derivation. If the derivation is empty, then \(f(...) = s = u = g(...)\) and therefore trivially \(f \rightarrowtail _{\mathcal {R}}^* g\). Otherwise, the derivation is of the shape Open image in new window for some nonvariable term \(t = h(...)\) and we obtain the inner induction hypothesis that \(f \rightarrowtail _{\mathcal {R}}^* h\). It remains to show \(h \rightarrowtail _{\mathcal {R}}^* g\) in order to conclude the proof. To this end, consider the step Open image in new window for some context C, substitution \(\sigma \), and rule Open image in new window such that Open image in new window . Now, by IH, we obtain that Open image in new window or Open image in new window or \(\mathsf {root}(s') \rightarrowtail _{\mathcal {R}}^* \mathsf {root}(t')\) for all Open image in new window . Thus, by Definition 11, we obtain that \(\mathsf {root}(l\sigma ) \rightarrowtail _{\mathcal {R}}\mathsf {root}(r\sigma )\). We conclude by a case analysis on the structure of the context C. If C is empty, that is Open image in new window , then \(h = \mathsf {root}(l\sigma ) \rightarrowtail _{\mathcal {R}}^* \mathsf {root}(r\sigma ) = g\) and we are done. Otherwise, \(h = \mathsf {root}(t) = \mathsf {root}(u) = g\) and therefore trivially \(h \rightarrowtail _{\mathcal {R}}^* g\). \(\square \)
Corollary 4
If \(f \rightarrowtail _{\mathcal {R}}^* g\) does not hold, then Open image in new window .
5.3 LookAhead Reachability in the Presence of Conditions
In the following definition we extend our lookahead technique from plain rewriting to conditional rewriting.
Definition 12
And we obtain a result similar to Theorem 2.
Lemma 7
If Open image in new window , then Open image in new window .
Example 12
6 Assessment
Experimental results for dependency graph analysis (TRSs).
Lookahead  

\(\mathsf {L}_{\mathcal {R}}^{0}\)  \(\mathsf {L}_{\mathcal {R}}^{1}\)  \(\mathsf {L}_{\mathcal {R}}^{2}\)  \(\mathsf {L}_{\mathcal {R}}^{3}\)  \(\mathsf {L}_{\mathcal {R}}^{8}\)  
None  UNSAT  0  104 050  105 574  105 875  105 993 
time (s)  33.96  38.98  38.13  39.15  116.52  
Corollary 1  UNSAT  307 207  328 216  328 430  328 499  328 636 
time (s)  38.50  42.71  42.72  43.00  66.82 
TRS Termination. For plain rewriting, we take all the 1498 TRSs from the TRS standard category of the termination problem data base version 10.6,^{6} the benchmark used in the annual Termination Competition [8], and overapproximate their dependency graphs. This results in 1 133 963 reachability constraints, which we call “edges” here. Many of these edges are actually satisfiable, but we do not know the exact number (the problem is undecidable in general).
For checking unsatisfiability of edges, we combine Corollary 2 for various values of k (0, 1, 2, 3, and 8), and either Corollary 1 or ‘None’. Here ‘None’ concludes unsatisfiability only for constraints that are logically equivalent to Open image in new window . In Table 1 we give the number of edges that could be shown unsatisfiable. Here, the ‘UNSAT’ row indicates the number of detected unsatisfiable edges and the ‘time’ row indicates the total runtime in seconds. (We ran our experiments on an Amazon EC2 instance model c5.xlarge: 4 virtual 3.0 GHz Intel Xeon Platinum CPUs on 8 GB of memory).
The starting point is \(\mathsf {L}_{\mathcal {R}}^{1}\) + None, which corresponds to the \(\mathsf {tcap}_{}\) technique, the method that was already implemented in NaTT before. The benefit of symbol transition graphs turns out to be quite significant, while the overhead in runtime seems acceptable. Moreover, increasing k of the lookahead reasonably improves the power of unsatisfiability checks, both with and without the symbol transition graph technique. In terms of the overall termination proving power, NaTT using only \(\mathsf {tcap}_{}\) solves 1039 out of the 1498 termination problems, while using \(\mathsf {L}_{\mathcal {R}}^{8}\) and Corollary 1, it proves termination of 18 additional problems.
CTRS Confluence. For conditional rewriting, we take the 148 oriented CTRSs of Cops,^{7} a benchmark of confluence problems used in the annual Confluence Competition [1]. Compared to version 1.5 of ConCon (the winner of the CTRS category in the last competition in 2018) our new version (1.7) can solve five more systems (that is a gain of roughly 3%) by incorporating a combination of Theorem 3, inductive symbol transition graphs (Corollary 4), and kfold lookahead (Lemma 7), where for the latter we fixed \(k = 1\) since we additionally have to control the level of conditional rewriting.
7 Related Work
Reachability is a classical topic in term rewriting; cf. Genet [7] for a survey. Some modern techniques include the treeautomatacompletion approach [5, 6] and a KnuthBendix completionlike approach [4]. Compared to these lines of work, first of all our interest is not directly in reachability problems but their (un)satisfiability. Middeldorp [12] proposed treeautomata techniques to approximate dependency graphs and made a theoretical comparison to an early termcapunifiability method [2], a predecessor of the \(\mathsf {tcap}_{}\)based method. It is indeed possible (after some approximations of input TRSs) to encode our satisfiability problems into reachability problems between regular tree languages. However, our main motivation is to efficiently test reachability when analyzing other properties like termination and confluence. In that setting, constructing tree automata often leads to excessive overhead.
Our work is inspired by the work of Lucas and Gutiérrez [11]. Their feasibility sequences serve the same purpose as our reachability constraints, but are limited to atoms and conjunctions. Our formulation, allowing other constructions of logic formulas, is essential for introducing lookahead reachability.
8 Conclusion
We introduced reachability constraints and their satisfiability problem. Such problems appear in termination and confluence analysis of plain and conditional rewriting. Moreover, we proposed two efficient techniques to prove (un)satisfiability of reachability constraints, first for plain and then for conditional rewriting. Finally, we implemented these techniques in the termination prover NaTT and the confluence prover ConCon, and experimentally verified their significance.
Footnotes
 1.
While in general we allow an arbitrary firstorder logical structure for formulas, for the purpose of this paper, negation and universal quantification are not required.
 2.
It is also possible to give a modeltheoretic account for these notions. However, the required preliminaries are outside the scope of this paper.
 3.
 4.
 5.
 6.
 7.
Notes
Acknowledgments
We thank Aart Middeldorp and the anonymous reviewers for their insightful comments. This work is supported by the Austrian Science Fund (FWF) project P27502 and ERATO HASUO Metamathematics for Systems Design Project (No. JPMJER1603), JST.
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