A Decision Procedure for Satisfiability in Separation Logic with

A Decision Procedure for Satisfiability in
Separation Logic with Inductive Predicates
James Brotherston ∗
Carsten Fuhs †
Juan A. Navarro P´erez ‡
University College London
We show that the satisfiability problem for the “symbolic heap”
fragment of separation logic with general inductively defined predicates — which includes most fragments employed in program verification — is decidable. Our decision procedure is based on the
computation of a certain fixed point from the definition of an inductive predicate, called its “base”, that exactly characterises its
A complexity analysis of our decision procedure shows that it
runs, in the worst case, in exponential time. In fact, we show that
the satisfiability problem for our inductive predicates is EXPTIMEcomplete, and becomes NP-complete when the maximum arity
over all predicates is bounded by a constant.
Finally, we provide an implementation of our decision procedure, and analyse its performance both on a synthetically generated set of test formulas, and on a second test set harvested from
the separation logic literature. For the large majority of these test
cases, our tool reports times in the low milliseconds.
Categories and Subject Descriptors F.3.1 [Logics and Meanings of Programs]: Specifying and Verifying and Reasoning about
Programs—Mechanical verification, assertions; F.2.2 [Analysis of
Algorithms and Problem Complexity]: Nonnumerical Algorithms
and Problems—Complexity of proof procedures
General Terms Algorithms, Theory, Verification
Keywords separation logic, inductive predicates, satisfiability, decision procedure
Separation logic [26] is an established and fairly popular formalism for verifying imperative, heap-manipulating programs. At the
∗ Research
supported by an EPSRC Career Acceleration Fellowship.
supported by EPSRC grants EP/K040863/1 and EP/H008373/2.
‡ Research supported by EPSRC grants EP/K040863/1 and EP/K032542/1.
† Research
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CSL-LICS 2014, July 14–18, 2014, Vienna, Austria.
Copyright is held by the owner/author(s).
ACM 978-1-4503-2886-9/14/07.
Nikos Gorogiannis
Middlesex University London
[email protected]
time of writing, there are a number of automatic program verification tools based on separation logic, such as SLAYER [5] and
A BDUCTOR [12], capable of establishing memory safety properties
of code bases extending into millions of lines [28]. These verification tools are highly dependent on the use of inductively defined
predicates to describe the shape of data structures held in memory, such as linked lists or trees. Currently, such predicates must be
hard-coded into these verification tools, which limits the range of
data structures that they can handle automatically. Thus, the next
step in automation is to handle general inductive predicates, which
might be provided to the analysis by the user, or even inferred automatically [10]. When one considers arbitrary inductive predicates,
however, it becomes much more difficult to determine whether a
given formula is satisfiable or not. This may lead to performance
degradation when time is spent on the analysis of inconsistent scenarios.
In this paper we address the latter problem by showing that the
satisfiability problem for the most commonly considered symbolic
heap fragment of separation logic, extended with general inductive
predicates, is in fact decidable. Our decision procedure rests upon
the observation that the satisfiability of each inductive predicate
can be precisely characterised by an approximation of its set of
models, an object which we refer to as the base of the predicate.
Roughly speaking, the base of a predicate records, for each distinct
way of constructing a satisfying model, the subset of arguments
of the predicate that are required to be allocated and distinct in
memory, as well as the equalities and disequalities that must hold
between these arguments. Since there are clearly only finitely many
possible such subsets and equality/disequality relations between
predicate arguments, the base can be straightforwardly computed
in finite time. Having computed the base of all required predicates,
deciding the satisfiability of arbitrary formulas in the fragment then
becomes a straightforward matter.
A complexity analysis of our decision procedure shows that,
in the worst case, it runs in exponential time (in the size of the
underlying set of inductive definitions). Essentially, this is because
our inductive definition schema allows us to construct predicates
that admit an exponential number of elements in their base. Indeed,
we show that the satisfiability problem for our inductive predicates
is EXPTIME-complete. Additionally, if the maximum arity of all
inductive predicates is bounded, then the satisfiability problem
becomes NP-complete.
We also provide an implementation of our decision procedure,
which is available online [1], and evaluate its performance over
three test sets: (a) a collection of example formulas drawn from
the literature on separation logic verification; (b) a large set of inductive predicates automatically generated by the C ABER predicate inference tool [10]; and (c) a set of synthetically generated
examples, varying in parameters such as the number of arguments
or recursive calls. Although exponential time performance can be
produced by a suitable choice of parameters, we find that for the
vast majority of examples, the algorithm terminates in a matter of
milliseconds. Consequently, we believe our decision procedure is
highly suitable for use as a black-box satisfiability checker in automated separation logic verification. We propose a number of direct
applications of such a satisfiability checker at the end of this paper
(in Section 7).
The problem of satisfiability for separation logic with inductive definitions was also recently considered by Iosif et al. [20].
Their paper establishes decidability results for both satisfiability
and entailment problems via an embedding into monadic second
order logic, but for technical reasons their results are restricted to a
fragment of the logic from which many natural data structure definitions are excluded; for example, inductive predicates may not
define structures with dangling data pointers (cf. Section 6.1). In
contrast, the fragment of separation logic for which we decide satisfiability includes a much larger class of inductive predicates, but
we do not consider the entailment problem for this fragment, which
is undecidable [2]. We also provide complexity results on the satisfiability problem for our fragment, and an implementation of our
decision procedure.
The remainder of this paper is structured as follows. We introduce our fragment of separation logic with inductive definitions in
Section 2, and then present our decision procedure and the proof
of its correctness in Section 3. Section 4 contains our complexity
analysis of our algorithm and of the satisfiability problem itself.
Section 5 describes the implementation of our decision procedure
and its evaluation. Section 6 surveys related work in the area, and
Section 7 concludes.
Inductive definitions in separation logic
Here we present our fragment of inductive definitions in separation
logic, following the approach in [9].
First, some preliminaries. A term is either a variable drawn from
the infinite set Var, or the constant symbol nil. We write Term for
the set of all terms. We also assume a fixed finite set P1 , . . . , Pn
of predicate symbols, each with an associated arity. We often write
vector notation to abbreviate tuples; e.g. we typically write x rather
than (x1 , . . . , xm ). We write πi (−) for the i-th projection function
on tuples, and sometimes abuse notation slightly by writing x ∈ x
to mean that x occurs in the tuple x. Finally, we write Pow(X) for
the powerset of a set X.
Definition 2.1. Spatial formulas F and pure formulas G are given
by the following grammar:
heap, Pi is a predicate symbol of arity ai , and x is a tuple of ai
distinct variables.
Variables in the body, Π : F , but not the head, Pi (x), of a rule
are implicitly existentially quantified, while multiple bodies of the
same head are interpreted as a disjunction. The formal semantics is
given in the next section.
Our strict formatting of the heads of inductive rules, with variable repetitions and occurrences of nil disallowed, is for technical convenience. It does not restrict expressivity since we can
achieve the same effect by placing equalities in the bodies of inductive rules. For example, a rule of the form ⇒ P (x, x, nil)
is not allowed by our schema, but the equivalent inductive rule
x = y, z = nil ⇒ P (x, y, z) is allowed.
We use a typical RAM model employing heaps of records. An
infinite set Val of values is assumed, of which an infinite subset
Loc ⊂ Val are locations, i.e., the addresses of heap cells. We also
assume a “nullary” value nil ∈ Val\Loc which is not the address of
any heap cell. A stack is a function s : Var → Val; we extend stacks
to terms by setting s(nil) =def nil , and extend stacks pointwise to
act on tuples of terms. We write s[x 7→ v] for the stack defined as
s except that (s[x 7→ v])(x) = v.
A heap is a partial function h : Loc *fin (Val List) mapping
finitely many locations to (arbitrary-length) tuples of values; we
dom(h) =def {` ∈ Loc | h(`) is defined}
and e to be the empty heap that is undefined everywhere. We
write ◦ to denote composition of heaps: if h1 and h2 are heaps,
then h1 ◦ h2 is the union of (partial functions) h1 and h2 when
dom(h1 ) ∩ dom(h2 ) = ∅, and undefined otherwise. We write
h[` 7→ v] for the heap defined as h except that (h[` 7→ v])(`) = v,
and just [` 7→ v] as a shorthand for e[` 7→ v]. We write Heap for
the set of all heaps.
Given an inductive rule set Φ, the relation s, h |=Φ F for
satisfaction of a pure or spatial formula F by the stack s and heap h
is defined as follows:
s, h |=Φ t1 = t2
s, h |=Φ t1 =
6 t2
s(t1 ) = s(t2 )
s(t1 ) 6= s(t2 )
s, h |=Φ emp
s, h |=Φ t 7→ t
dom(h) = {s(t)} and h(s(t)) = s(t)
s, h |=Φ Pi t
s, h |=Φ F1 ∗ F2
(s(t), h) ∈ JPi KΦ
h = h1 ◦ h2 and s, h1 |=Φ F1
and s, h2 |=Φ F2
where t ranges over terms, Pi over the predicate symbols and t over
tuples of terms (matching the arity of Pi in Pi t). We write PureSet
for the set of all finite sets of pure formulas.
A symbolic heap is given by Π : F , where F is a spatial formula
and Π ∈ PureSet (note that Π should be read intuitively as the
conjunction of its elements). Whenever one of Π, F is empty, we
will omit the colon.
We write the substitution notation F [t/x] for the result of simultaneously replacing all occurrences of the variable x by the term t
in the formula F . Substitution extends to sets of formulas in the
obvious way.
where the semantics JPi KΦ of the inductive predicate Pi under
Φ is defined below. If Π : F is a symbolic heap then we write
s, h |=Φ Π : F to mean that s, h |=Φ F and s, h |=Φ G for all
G ∈ Π; equivalently, we say that (s, h) is a model of Π : F (w.r.t.
Φ). A symbolic heap Π : F is satisfiable (w.r.t. Φ) if it has at least
one model.
We remark that satisfaction of pure formulas does not depend
on the heap nor on the inductive rules: we write s |= Π, where
Π ∈ PureSet, to mean that s, h |=Φ Π for any heap h and inductive
definition set Φ.
The following definition gives the standard semantics of the
inductive predicate symbols P according to a fixed inductive rule
set Φ, i.e., as the least fixed point of an n-ary monotone operator
constructed from Φ:
Definition 2.2. An inductive rule set is a finite set of inductive
rules, each of the form Π : F ⇒ Pi x, where Π : F is a symbolic
Definition 2.3. First, for each predicate Pi ∈ P with arity ai say,
we define τi = Pow(Val ai × Heap). Furthermore, we partition
F ::= emp | t 7→ t | Pi t | F ∗ F
G ::= t = t | t 6= t
the rule set Φ into Φ1 , . . . , Φn , where Φi is the set of all inductive
rules in Φ of the form Π : F ⇒ Pi x.
We let each Φi be indexed by j (i.e., Φi,j is the j-th rule defining
Pi ), and for each inductive rule Φi,j of the form Π : F ⇒ Pi x, we
define the operator ϕi,j : τ1 × . . . × τn → τi by:
ϕi,j (Y) =def {(s(x), h) | s, h |=X Π : F }
where Y ∈ τ1 ×. . .×τn and |=Y is the satisfaction relation defined
above, except that JPi KY =def πi (Y). We then finally define the
tuple JPKΦ ∈ τ1 × . . . × τn by:
JPKΦ =def µY. ( j ϕ1,j (Y), . . . , j ϕn,j (Y))
We write JPi KΦ as an abbreviation for πi (JPKΦ ).
As in [9], Definition 2.3 interprets the inductive rules Φi as
exhaustive, disjunctive clauses of the definition of the predicate Pi .
A decision procedure for satisfiability of
inductive predicates
In this section, we present our decision procedure for satisfiability
in the symbolic heap fragment of separation logic with inductive
predicates, given in Section 2.
Our first definition is used to normalise terms in pure formulas
with respect to a given set of variables. This is then used in our
second definition, which provides a satisfiability-preserving means
of restricting the variables which occur in a set of pure formulas.
Definition 3.1. Let Π ∈ PureSet and let x be a tuple of distinct
variables. The equalities in Π determine an equivalence relation
among terms. We say that an equivalence class is observable (from
x with respect to Π), if the class contains some term in x ∪
{nil}. A term is observable just in case it belongs to an observable
equivalence class.
For each equivalence class, we choose a term from the class,
called its canonical representative. For observable classes, we
choose nil to be the representative if nil is in the class, otherwise we
choose a variable from x. For non-observable classes, the choice is
arbitary. The normal form of a term t, denoted htiΠ,x , is defined to
be t itself when t ∈ x ∪ {nil}, or the canonical representative of its
equivalence class otherwise.
Definition 3.2. Let Π ∈ PureSet and let x be a tuple of distinct
variables. We denote by Π x the formula obtained after: (i)
removing all pure formulas, t1 = t2 or t1 6= t2 , where at least
one of t1 or t2 is not observable; and (ii) replacing all remaining
terms t by their normal forms htiΠ,x .
Note that Definition 3.2 ensures every term occurring in Π x is
either nil, or a variable from x. The satisfiability-preserving nature
of this variable restriction is formalised by the following pair of
Lemma 3.3. Let Π be a finite set of pure formulas. If s |= Π and
s(x) = s0 (x) then s0 |= Π x.
Proof. Since s |= Π, we know that s must assign the same value to
all variables sharing the same equivalence class. Thus s |= Π x,
since both dropping formulas from Π and replacing terms by their
normal forms preserve its satisfiability.
Now, Π x is restricted to contain only variables from x and,
since s and s0 agree on those variables, s0 |= Π x.
Lemma 3.4. Let Π be a finite set of pure formulas. If Π is satisfiable and s |= Π x then there exists a stack s0 with s0 (x) = s(x)
such that s0 |= Π.
Furthermore, given any finite set W ⊆ Loc, we can choose s0
such that s0 (y) ∈
/ W for all non-observable variables y.
define base Φ P:
Y := (λt1 . ∅, . . . , λtn . ∅)
repeat until Y reaches a fixed point:
pick a rule Φi,j ∈ Φ with head Pi x
for each Pj` (x` ) in the body of Φi,j :
pick a base pair (V` , Π` ) ∈ Yj` (x` )
take y1 , . . . , yk from all y` 7→ u` in the body of Φi,j
take Π0 the pure part in the body of Φi,j
V := V1 ∪ · · · ∪ Vm ∪ {y1 , . . . , yk }
Π :=
V ∪ Π0 ∪ Π1 ∪ · · · ∪ Πm
if Π is satisfiable:
add alloced (V, x, Π)[t/x], (Π x)[t/x] to Yi (t)
return Y
Figure 1. Pseudocode for the computation of base Φ P in Defn. 3.5.
Proof. Let Y1 , . . . , Yn be the equivalence classes of all variables occurring in, and determined by, Π. Choose distinct values
`1 , . . . , `n in Loc such that each `i ∈
/ s(x ∪ {nil}) ∪ W . This is
possible because the number of locations in Loc is infinite, while
the number of forbidden locations is finite. We then define
s(hxiΠ,x ) if x is observable
s (x) =
otherwise, where x ∈ Yi .
By construction s and s0 agree on variables from x, i.e. directly
observable ones, and s0 (y) ∈
/ W when y is not observable. It is
easy to check now that s0 satisfies all pure formulas in Π.
• For every t1 = t2 ∈ Π: By definition, t1 and t2 are in the same
equivalence class. Thus if t1 is observable then so, necessarily,
is t2 , and furthermore the formula ht1 iΠ,x = ht2 iΠ,x is in
Π x. Therefore, by assumption, s |= ht1 iΠ,x = ht2 iΠ,x ,
and since s0 (ti ) = s(hti iΠ,x ), we have s0 |= t1 = t2 .
Otherwise, if t1 is not observable, then neither is t2 ; but, since
they share an equivalence class, we have s0 (t1 ) = `i = s0 (t2 )
for some i by construction, and so s0 |= t1 = t2 as required.
• For every t1 6= t2 ∈ Π: If both t1 and t2 are observable,
then ht1 iΠ,x 6= ht2 iΠ,x is in Π x, so s |= ht1 iΠ,x 6=
ht2 iΠ,x by assumption. Since s0 (ti ) = s(hti iΠ,x ), it follows
that s0 |= t1 6= t2 .
If one of the terms, say t1 , is observable but the other, t2 ,
is not, then s0 (t2 ) was explicitly chosen to be different to
s(ht1 iΠ,x ) = s0 (t1 ), and so s0 |= t1 6= t2 .
If neither t1 nor t2 is observable, then they must be in different
equivalence classes, otherwise Π would be unsatisfiable, contrary to assumption. Thus, by construction, s0 |= t1 6= t2 .
We now present our main definition, which explains how to extract from a fixed inductive rule set Φ all the information needed
to decide the satisfiability of any symbolic heap over Φ. The pseudocode in Figure 1 provides an informal aid to navigate the steps
of the computation.
Definition 3.5 (Base pair computation). First, for each predicate
symbol Pi with associated arity ai , where 1 ≤ i ≤ n, we define
σi =def Termai → Pow(MPow(Var) × PureSet)
where MPow(Var) denotes the set of allN
multisets of variables. For
any finite V ∈ MPow(Var) we define
V ∈ PureSet to be the
set of pure formulas comprising:
• all formulas of the form x 6= x0 such that x and x0 are different
N of V (note this means that if V contains duplicates
V is unsatisfiable);
• the formula x 6= nil for every element x of V .
For any finite multiset V of variables, any tuple of variables
x, and any Π ∈ PureSet, let alloced (V, x, Π) denote the multiset containing the term hyiΠ,x for each y N∈ V observable
from x with respect to Π. Note that if Π ∪
V is satisfiable,
then alloced (V, x, Π) contains only variables (no nil) and no duplicates. In this way we represent, using only variables from x, the
observable locations allocated via V .
The inductive rule set Φ is partitioned into Φ1 , . . . , Φn with
each Φi further indexed by j as in Definition 2.3. Without loss of
generality, we consider each Φi,j ∈ Φ to be written in the form
Π0 : y1 7→ u1 ∗ . . . ∗ yk 7→ uk ∗
Pj1 (x1 ) ∗ . . . ∗ Pjm (xm ) ⇒ Pi x,
where Π0 ∈ PureSet. We use the inductive rule Φi,j to define
an operator Ψi,j : σ1 × . . . × σn → σi as follows. If Y =
(Y1 , . . . , Yn ) where each Yi ∈ σi , and t is a tuple of ai terms,
then Ψi,j (Y) : σi sends t to the set of pairs (we shall call these
pairs also base pairs)
alloced (V, x, Π)[t/x], (Π x)[t/x]
that satisfy the following:
V = V1 ∪ · · · ∪ Vm ∪ {y1 , . . . , yk },
V ∪ Π0 ∪ Π1 ∪ · · · ∪ Πm is satisfiable,
and ∀1 ≤ ` ≤ m. (V` , Π` ) ∈ Yj` (x` ).
Note that the substitution [t/x] in the above is defined pointwise
over tuples; this is well defined since x is a tuple of distinct variables, as per Definition 2.2. Note also that while alloced (V, x, Π)
might be viewed as a set (i.e., without repetitions), it is important
to understand alloced (V, x, Π)[t/x] as a multiset since the substitution may map different variables to the same term.
We then define base Φ P ∈ σ1 × . . . × σn as follows:
base Φ P =def µY. ( j Ψ1,j (Y), . . . , j Ψn,j (Y))
where,Sby pointwise extension of the union to act on functions, we
write j Ψi,j (Y) to denote the function mapping a tuple of terms
t to the set j Ψi,j (Y)(t). We also write base Φ Pi to abbreviate
πi (base Φ P).
Example 3.6. Consider the following set Φ of inductive rules:
Φ1,1 :
x = nil ⇒ P (x)
Φ1,2 :
Φ2,1 :
x 6= nil : Q(x, x) ⇒ P (x)
y = nil, x 6= nil : x 7→ (d, c) ∗ P (d) ⇒ Q(x, y)
Φ2,2 :
y 6= nil : y 7→ (d, c) ∗ Q(x, c) ⇒ Q(x, y).
These rules are an intermediate product of the analysis done by the
tool C ABER [10] on code that traverses a list of lists, and are found
to be ultimately unsatisfiable by our algorithm, triggering backtracking in C ABER. In accordance with Definition 3.5, we define:
σ1 = Term → Pow(MPow(Var) × PureSet)
σ2 = Term × Term → Pow(MPow(Var) × PureSet) .
That is, σ1 and σ2 are function spaces mapping suitably many terms
to sets of base pairs for P and Q, respectively. We can compute
base Φ (P, Q) by iteratively computing the sequence of its fixed
point approximants ((Y1i , Y2i ) ∈ σ1 × σ2 )i≥0 in the standard way.
That is, we begin with an empty base for both predicates, namely
Y0 = (λx. ∅, λx, y. ∅), and iteratively apply the operators Ψi,j
given by Definition 3.5.
First iteration: The inductive rule Φ1,1 for P (x) can be applied
to Y0 , yielding V = ∅ and Π = {x = nil}. The set Π is satisfiable
and, therefore,
Ψ1,1 (Y10 , Y20 ) = λx. (∅, {x = nil}) .
Since Y0 does not contain any base pairs, the remaining rules
cannot be applied to it, and so Y1 = (Y11 , Y21 ) where
Y11 = λx. (∅, {x = nil})
Y21 = λx, y. ∅ .
Second iteration: The inductive rule Φ1,1 generates the same
base pair as before, while Φ1,2 and Φ2,2 are still not applicable to
Y1 . However, the rule Φ2,1 defining Q(x, y) can be now applied,
taking (∅, {x = nil}) from the current base Y11 of P (x). After
restricting and normalising to the set of observable variables, this
produces the new base pair ({x}, {y = nil, x 6= nil}). Thus, at the
end of the iteration, Y2 = (Y12 , Y22 ) where
Y12 = λx. (∅, {x = nil})
Y22 = λx, y. ({x}, {y = nil, x 6= nil}) .
Third iteration: The rule Φ1,2 is now applicable to Y2 , instantiating the current pair for Q(x, x), but yields an unsatisfiable Π. Similarly, the rule Φ2,2 can be applied, taking the base pair
({x}, {c = nil, x 6= nil}) from Y22 instantiated to Q(x, c). After
a suitable variable restriction and normalisation, this produces the
following new base pair for Q(x, y):
({x, y}, {x 6= y, y 6= nil, x 6= nil})
Thus Y = (Y13 , Y23 ) where
Y13 = λx. (∅, {x = nil})
({x}, {y = nil, x 6= nil}),
Y23 = λx, y.
({x, y}, {x 6= y, y 6= nil, x 6= nil})
On the fourth iteration no new base pairs are produced — both
Φ1,2 and Φ2,2 can be applied to the newly created base pair from
the latest iteration, but the former yields an unsatisfiable Π and the
latter reproduces the same base pair — so we have reached a fixed
point, and base Φ (P, Q) = (Y13 , Y23 ).
The next two lemmas formalise the fact that base Φ P (x) precisely characterises the satisfiability of P (x).
Lemma 3.7 (Soundness). Given a base pair (V, Π) ∈
(base Φ Pi )(t), a stack s such that s |= Π, and a finite set
W ⊆ Loc \ s(V ), there exists a heap h such that s, h |=Φ Pi t and
dom(h) ∩ W = ∅.
Proof. We proceed by fixed point induction on the definition of
base Φ P. That is, we assume the lemma holds for a tuple of functions Y = (Y1 , . .S. , Yn ) ∈ σ1 × · S
· · × σn , and we must show
it also holds for ( j Ψ1,j (Y), . . . , j Ψn,j (Y)). Thus, assume
(V, Π) ∈ Ψi,j (Y)(t) with s |= Π and s(V ) ∩ W = ∅. We require
to find an h with s, h |= Pi t and dom(h) ∩ W = ∅.
By assumption, there is a rule Φi,j of the form (IndRule) with
(V, Π) ∈ Ψi,j (Y)(t). Therefore V = alloced (V 0 , x, Π0 )[t/x]
and Π = (Π0 x)[t/x], where the following hold:
V 0 = V1 ∪ · · · ∪ Vm ∪ {y1 , . . . , yk },
N 0
Π0 =
V ∪ Π0 ∪ Π1 ∪ · · · ∪ Πm is satisfiable,
∀1 ≤ ` ≤ m. (V` , Π` ) ∈ Yj` (x` ).
By the lemma assumption and usual substitution facts, we have that
s[x 7→ s(t)] |= Π0 x and
s[x 7→ s(t)](alloced (V 0 , x, Π0 )) ∩ W = ∅.
Since Π0 is satisfiable, we can apply Lemma 3.4 to obtain s0 with
s0 (x) = s[x 7→ s(t)](x) = s(t) and s0 |= Π0 . Furthermore,
the lemma tells us that for any non-observable y ∈ V 0 , we have
/ W . Otherwise, if y ∈ V 0 is observable from x w.r.t. Π0 , there
is an x ∈ x such that y and x are in the same Π0 -equivalence
class. It follows that s0 (y) = s0 (x) = s[x 7→ s(t)](x) and, as
x ∈ alloced (V 0 , x, Π0 ) but s[x 7→ s(t)](alloced (V 0 , x, Π0 )) ∩
W = ∅, it must be the case that s0 (y) ∈
/ W . Thus, for all y ∈ V 0
we have s (y) ∈
/ W , that is s (V ) ∩ W = ∅.
We now show that there are heaps h1 , . . . , hm such that, for all
1 ≤ ` ≤ m, we have both s0 , h` |=Φ Pj` x` and dom(h` ) ∩ W` =
∅, where W` is defined as follows:
[ 0
s (Vq ) .
W` =def W ∪ s0 ({y1 , . . . , yk }) ∪
dom(hp ) ∪
To do this, we inductively assume that we have constructed the
chain of heaps (hp )1≤p<` , and show how to construct h` .
Inductive construction of h` from (hp )1≤p<` : We claim and prove
that s0 (V` ) ∩ W` = ∅. We have shown that s0 (V 0 ) ∩ W = ∅; since
V` ⊆ V 0 , it follows that s0 (V` ) ∩ W = ∅. Next, let p < ` and
note that the induction hypothesis implies dom(hp ) ∩ Wp = ∅,
which in turn implies
S dom(hp ) ∩ s (V` ) = ∅ by definition of
Wp . Thus s (V` ) ∩ p<` dom(hp ) = ∅. Next, let q > ` and
N 0
notice that s0 (V` ) ∩ s0 (Vq ) = ∅ because s0 |=
V , which
guarantees that s S
is injective on V ⊇ V` , Vq . It therefore follows that s0 (V` ) ∩ `<q≤m s0 (Vq ) = ∅. Also, for a similar reason,
s0 (V` ) ∩ s0 ({y1 , . . . , yk }) = ∅.
Now (V` , Π` ) ∈ Yj` (x` ), where s0 |= Π` and s0 (V` )∩W` = ∅,
so by the main induction hypothesis (of the fixed point induction for
the present lemma) there is a heap h` such that s0 , h` |=Φ Pj` x`
and dom(h` ) ∩ W` = ∅. This completes the construction of h` .
We continue with the main proof. Note that h1 ◦ . . . ◦ hm
is defined because dom(h` ) ∩ W` = ∅ for all 1 ≤ ` ≤ m
implies that dom(h1 ), . . . , dom(hm ) are all disjoint from each
other. Now we define a heap h0 whose domain is s0 ({y1 , . . . , yk })
as follows: h0 (s0 (yi )) =def s0 (ui ) for each 1N
≤ i ≤ k. Note that
h0 ◦ (h1 ◦ . . . ◦ hm ) is defined because s0 |=
V 0 ensures that s0
is injective on {y1 , . . . , yk } ⊆ V 0 and dom(h` ) ∩ W` = ∅ ensures
that dom(h` ) ∩ dom(h0 ) = ∅ for each 1 ≤ ` ≤ m. Thus, defining
h =def h0 ◦ h1 ◦ . . . ◦ hm , we obtain:
s0 , h |=Φ Π0 : y1 7→ u1 ∗ . . . ∗ yk 7→ uk ∗ Pj1 x1 ∗ . . . ∗ Pjm xm
which implies that s , h |=Φ Pi x (by applying the operator ϕi,j ).
Thus, since s0 (x) = s(t), we obtain s, h |=Φ Pi t.
We have dom(h) ∩ W = ∅ as required because
dom(h` ) ∩ W ⊆ dom(h` ) ∩ W` = ∅,
for each 1 ≤ ` ≤ m, and dom(h0 ) = s0 ({y1 , . . . , yk }) ⊆ s0 (V 0 )
while s0 (V 0 ) ∩ W = ∅.
Lemma 3.8 (Completeness). If s, h |=Φ Pi t, there is a base pair
(V, Π) ∈ (base Φ Pi )(t) such that s(V ) ⊆ dom(h) and s |= Π.
Proof. We have that (s(t), h) ∈ JPi KΦ , and apply fixed point
induction on the definition of JPKΦ . That is, we assume the
lemma holds for X = (X
S1 , . . . , Xn ) ∈ τS1 × . . . × τn and
must show that it holds for ( j ϕ1,j (X), . . . , j ϕn,j (X)). Thus,
assuming that (s(t), h) ∈ ϕi,j (X), we must then find a pair
(V, Π) ∈ (base Φ Pi )(t) with s(V ) ⊆ dom(h) and s |= Π.
By assumption, there is an inductive rule Φi,j of the form
(IndRule) such that (s(t), h) ∈ ϕi,j (X). By construction this
means that s(t) = s0 (x) for some stack s0 , and we have
s0 , h |=X Π0 : y1 7→ u1 ∗ . . . ∗ yk 7→ uk ∗ Pj1 x1 ∗ . . . Pjm xm
Thus s0 |= Π0 and h = h1 ◦ . . . ◦ hk ◦ h01 ◦ . . . ◦ h0m , where
s0 , h` |=Φ y` 7→ u` for all 1 ≤ ` ≤ k, and (s0 (x` ), h0` ) ∈ Xj`
for all 1 ≤ ` ≤ m. Then, by the induction hypothesis, for all
1 ≤ ` ≤ m there are pairs (V` , Π` ) ∈ (base Φ Pj` )(x` ), such that
s0 (V` ) ⊆ dom(h0` ) and s0 |= Π` . Now we define
V` ∪ {y1 , . . . , yk } .
V =def
As dom(h1 ), . . . , dom(hk ), dom(h01 ), . . . , dom(h0m ) are all disjoint, where dom(h` ) = {s0 (y` )} for each 1 ≤ ` ≤ kNand
dom(h0` ) ⊇ s0 (V` ) for each 1 ≤ ` ≤ m, we have s0 |=
(i.e., h would be undefined if s0 were not injective on V ). Putting
everything together, we have
s0 |= Π0 ∪
V ∪ Π1 ∪ · · · ∪ Πm and s0 (V ) ⊆ dom(h) .
Thus Π = Π0 ∪
V ∪ Π1 ∪ · · · ∪ Πm is satisfiable, and so
(V ∩ x)[t/x], (Π x)[t/x] ∈ (base Φ Pi )(t)
where the substitution [t/x] is well defined as x is a tuple of distinct
variables. From the fact that s0 (x) = s(t), it also follows that
s0 (x) = s[x 7→ s(t)](x). By Lemma 3.3, this gives us
s[x 7→ s(t)] |= Π x and s[x 7→ s(t)](V ∩ x) ⊆ dom(h) .
Thus, from usual facts about substitution, we obtain
s |= (Π x)[t/x] and s(V ∩ x)[t/x] ⊆ dom(h) .
Thus there exists (V, Π) ∈ (base Φ Pi )(t) such that s |= Π and
s(V ) ⊆ dom(h) as required.
We are now positioned to state our main decidability results.
Theorem 3.9. A formula Pi t is satisfiable w.r.t. an inductive rule
set Φ iff there is a base pair (V, Π) ∈ (base Φ Pi )(t) such that Π is
Furthermore, for any given Φ, the fixed point computation of
base Φ P terminates after a finite number of steps.
Thus it is decidable whether a formula of the form Pi t is
satisfiable w.r.t. Φ.
Proof. Regarding correctness, the “if” direction is immediate from
Lemma 3.7 by taking W = ∅; the “only if” direction follows from
Lemma 3.8.
For termination, each predicate Pi has a finite arity ai and, since
each of the pairs (V, Π) ∈ base Φ Pi (t) may only use terms from
the finitely many in t, the final total number of pairs in base Φ Pi (t)
is bounded by a finite constant. All of the functions in the tuple
base Φ P, mapping terms to base pairs, are thus finitely described in
a straightforward way.
Corollary 3.10. It is decidable whether an arbitrary symbolic heap
is satisfiable w.r.t. an inductive rule set Φ.
Proof. Let Π : F be a symbolic heap, and let Ψ be the inductive
rule set Φ ∪ {Π : F ⇒ Q}, where Q is a new 0-ary predicate
not occurring in Φ or Π : F . Clearly the symbolic heap Π : F is
satisfiable w.r.t. Φ if and only if Q is satisfiable w.r.t. Ψ, which is a
decidable problem according to Theorem 3.9.
In this section, we investigate the complexity of the satisfiability
problem for inductive predicates in separation logic, which is addressed by the decision procedure in the previous section.
Definition 4.1. We define the decision problem PREDSAT as
having instances (Φ, P ) where Φ is an inductive rule set and P
is a predicate defined in Φ. The pair (Φ, P ) is a yes-instance of
PREDSAT if P x is satisfiable w.r.t. Φ, where x is an arbitrary
tuple of distinct variables.
For any k ∈ N, define k-PREDSAT to be PREDSAT restricted
to the situation where all predicates in Φ have arity ≤ k.
In the following, the length of the encoding of a mathematical object o (e.g., set, pair, formula, etc.) under some reasonable
(non-unary) encoding scheme will be denoted by kok. Clearly,
k(Φ, P )k = O(kΦk), as P can be represented by a pointer of
dlog2 |Φ|e bits. Finally we fix B =def {>, ⊥}.
Let α be the maximum arity of any predicates in Φ plus one;
clearly, α = O(kΦk) as the maximum arity must be of the order of
the length of the entire object. The length of a base pair is bounded
by α + 2α2 = O(kΦk2 ) (size of the variable set plus the size of
the longest pure formula), and the number of distinct base pairs for
any predicate in Φ is bounded by
N =def 2α+2α .
Finally L is the maximum number of predicate occurrences in any
rule; as such, L = O(kΦk).
Lemma 4.2. Let Φi,j be a rule in the form of (IndRule) and
let T = hB1 , . . . , Bm i be a tuple of base pairs for the formulas
Pj1 (x1 ), . . . , Pjm (xm ) respectively. Computing whether a base
pair B 0 can be generated from T and Φi,j via Definition 3.5 takes
time polynomial in kΦi,j k and kT k.
Proof. Immediate from Definition 3.5 and the fact that satisfiability
of conjunctions of pure formulas is in P (see, e.g., [3]).
Definition 4.3. Let Φ be a set of inductive rules. A rule instance
r = (B, i, j, r1 , . . . , rm ), for i, j, m ≥ 0 is such that: (a) the rule
Φi,j has m predicates in its body; (b) B is a base pair for Pi ;
(c) r1 , . . . , rm are rule instances; and (d) B can be produced by
running the body of the loop in Fig. 1 using Φi,j and the base pairs
associated with r1 , . . . , rm . We say that the rule instance r defined
as above supports predicate Pi .
In effect, a rule instance is a base pair together with the witnessing information for its production via our algorithm.
Lemma 4.4. A predicate P defined in Φ is satisfiable iff there is a
rule instance that supports P .
Proof sketch. It is simple to modify the algorithm in Definition 3.5
so that instead of base pairs, it generates rule instances. Note that
the rule instances within a rule instance r must be generated prior
to r and thus can be represented by pointers, therefore not altering
the space complexity of the algorithm.
Lemma 4.5. k-PREDSAT is in NP, for any k ∈ N.
Proof. We define a non-deterministic algorithm as follows. Let
(Φ, P ) be the input instance. As stated in the beginning of this section, there can be up to N distinct base pairs for each predicate in Φ.
Thus, for each predicate Pi in Φ we guess a set of up to N entries of
the form hB, i, j, E1 , . . . , Em i, where B is a base pair, i, j, m ≥ 0,
and E` is a pointer to an entry for 1 ≤ ` ≤ m. Each entry requires
O(kΦk3 ) time to generate: B takes O(α + 2α2 ) = O(kΦk2 )
time; pointers i, j take 2dlog2 |Φ|e = O(kΦk) time; pointers
E1 , . . . , Em take mdlog2 N e = O(L(α + 2α2 )) = O(kΦk3 )
time. By assumption, the arity of all predicates in Φ is bounded by
a constant k, thus N is also constant. This means the total number
of guesses is polynomial in the input size.
We now check that this set of entries represents a rule instance
supporting P . To do this we check: (a) that each entry is wellformed, i.e., that rule Φi,j involves m predicates, and that each
E` points to an entry for predicate Pj` ; and (b) that each entry E
can be generated through our algorithm using entries E1 , . . . , Em
and rule Φi,j . Then, we check that the set of entries represents a
directed acyclic graph when we view the pointers E` as edges. This
will guarantee that there are no cycles and that, therefore, the set
of entries represents a set of rule instances. These operations take
polynomial time in N and |Φ| by Lemma 4.2 and standard facts
from graph theory. Finally, we check for an entry (rule instance)
supporting P .
Since we can check the guessed entries in polynomial time, we
can decide satisfiability in non-deterministic polynomial time.
Lemma 4.6. PREDSAT is in EXPTIME.
Proof. Each step of the algorithm in Definition 3.5 picks a rule
from Φ and looks for a combination of base pairs that yields a new
base pair. Thus it may go through up to N L combinations of base
pairs. In the worst case all rules have to be scanned, checking up
to |Φ|N L combinations. If no new base pair can be generated then
the algorithm has reached a fixed point.
As argued above, there can be at most |Φ|N distinct base
pairs. In the worst case one pair is generated in each step. Thus
|Φ|N steps are required with total time |Φ|N |Φ|N L p(kΦk) =
O(2poly(kΦk) ) where p(kΦk) is the time taken to check a single
combination of base pairs according to Lemma 4.2.
Our lower bound results use standard facts about boolean circuits, with and without inputs. We make a few common simplifying assumptions: (a) inputs and constants are gates with no inputs
and one output; (b) every circuit includes exactly two constants, >
and ⊥; (c) there is one more kind of gate, the NAND gate (denoted
by ↑); (d) gates are ordered so that if gate i has inputs l, r, then
i > l and i > r; (e) inputs, >, ⊥ and NAND gates appear in this
order; (f) the output of the maximal gate is the output of the circuit.
We first define a mapping from circuits to inductive rule sets.
Definition 4.7. Let C be a boolean circuit with n gates and k
inputs. Thus, the constant gates are k + 1 and k + 2 (> and ⊥
respectively). Define a set of rules Ψ as follows.
x 6= nil ⇒ T (x)
x = nil ⇒ F (x)
F (x) ∗ T (z) ⇒ N (x, y, z)
Ψ =def
F (y) ∗ T (z) ⇒ N (x, y, z)
T (x) ∗ T (y) ∗ F (z) ⇒ N (x, y, z)
Also define the set of rules ΦC as follows.
) ∗ F (xk+2 ) ∗
 T (xk+1
∗ni=k+3 N (xli , xri , xi ) ⇒ P (xn )
ΦC =def
P (xn ) ∗ F (xn ) ⇒ Q⊥ 
Above, li (resp. ri ) denotes the left (right) input of gate i. Clearly,
kΦC k = O(kCk) and the maximum arity is 3.
We will often have to go from boolean tuples to stacks and back.
The following definition provides appropriate mappings.
Definition 4.8. Fix some τ ∈ Val such that τ 6= nil . Let b ∈ B,
B ∈ Bn and x ∈ Varn . Define functions sval : B → Val,
bval : Val → B, stack sxB and boolean n-tuple Bsx as
if b = >
> if v 6= nil
sval(b) =def
bval(v) =def
nil if b = ⊥
⊥ if v = nil
sxB (xi ) =def sval(Bi )
(Bsx )i =def bval(s(xi ))
where i = 1, . . . , n.
Theorem 4.9. k-PREDSAT is NP-complete for k ≥ 3.
Proof. Membership in NP follows from Lemma 4.5. Hardness
follows by reducing from CIRCUITSAT via Definition 4.7. We
show that there is a boolean k-tuple B such that C(B) = > iff
(ΦC , Q> ) ∈ k-PREDSAT, by proving that for any b there exists
B such that C(B) = b iff (ΦC , Qb ) ∈ k-PREDSAT.
(⇒) Suppose there exists B such that C(B) = b. Extend B to
the n-tuple B0 such that for all i = 1, . . . , n, Bi0 is equal to the
output of gate i. Clearly, B0 is well defined. Let sˆ =def sxB0 . It is
easy to see that if b = > = Bn0 then sˆ |=Ψ T (xn ) and that if
b = ⊥ then sˆ |=Ψ F (xn ). Thus in order to show that sˆ |=ΦC Qb
it remains to show that sˆ |=ΦC P (xn ). If n ≤ k + 2 this is
immediate as the output of C is either a constant or an input. If n is
a NAND gate then we must show that for every i = k + 3, . . . , n,
Bl0i ↑ Br0 i = Bi0 implies sˆ |=ΦC N (xli , xri , xi ), where li , ri are
the inputs of gate i. This holds by the definitions of predicate N
and stack sˆ.
(⇐) Suppose there is a stack s such that s |=ΦC Qb . We assume
ˆ =def Bsx . Set B as the kw.l.o.g. that s |=ΦC P (xn ) also. Let B
ˆ We use induction to prove that for every i = 1, . . . , n,
prefix of B.
ˆi . The
if the inputs are set to B then the output of gate i is B
case where i = 1, . . . , k + 2 is trivial. If i = k + 3, . . . , n
then it is a NAND gate and has inputs li , ri < i whose values
ˆl , B
ˆr . It is then simple to verify that if
we know are equal to B
ˆl ↑ B
ˆr = B
ˆi , by the
s |=Ψ N (xli , xri , xi ) then it follows that B
definitions of N and B.
If b = > then by assumption s |=ΦC P (xn ) ∗ T (xn ), so
ˆn = > and thus C(B) = >. The case b = ⊥ is similar.
For EXPTIME-hardness we will encode natural numbers as
boolean tuples as well as formulas, as follows.
Definition 4.10. We use iB to denote the boolean n-tuple encoding
an integer 1 ≤ i ≤ 2n . Note that, for convenience, we set 1B = ⊥n
and (2n )B = >n . In addition, we use a formula encoding in terms
of the predicates T and F as seen above.
T (x) if b = >
frmb (x) =def
F (x) if b = ⊥
boolB (x) =def frmB1 (x1 ) ∗ . . . ∗ frmBn (xn )
numi (x) =def bool(iB ) (x)
The circuit value problem (CVP) has as instances input-free
circuits. A circuit C is a yes-instance of CVP iff C() = >. The
succinct circuit value problem (SCVP) has as instances input-free
circuits generated by a pair of circuits as explained below. An
instance of SCVP is a yes-instance iff for the generated circuit
C, C() = >. Formally, an instance SC = (L, R) of SCVP
consists of two circuits, each of 2n inputs. In the represented
circuit C, for all i, j ∈ {1, . . . , 2n } such that L(iB , j B ) = >
(resp. R(iB , j B ) = >), it holds that i > j, i is a NAND gate and its
left (right) input is j. For simplicity we assume that circuits always
use 2n gates and that their output is that of gate 2n . Also recall that
for a k-input circuit, gates k +1 and k +2 are > and ⊥ respectively.
Definition 4.11. Let C be a k-input circuit with n gates. Then,
define ΦrC as follows, where Ψ is as in Definition 4.7.
 T (xk+1 ) ∗ F (xk+2 ) ∗ T (xn ) ∗
ΦrC =def Ψ ∪
 ∗n
i=k+3 N (xli , xri , xi ) ⇒ P (x1 , . . . , xk )
This definition turns the circuit C into a relation P over the variables representing the inputs of C. This will simplify presentation
when we are only interested for inputs that make the output equal
to >. In particular, the following lemma holds.
Lemma 4.12. For any circuit C on k-inputs and any B ∈ Bk ,
C(B) = > iff s |=ΦrC P (x) ∗ boolB (x) for some stack s.
Proof. Analogous to the proof of Theorem 4.9.
Lemma 4.13. For any k-input circuit C and stacks s, s0 , if
Bsx = Bsx then s |=ΦrC P (x) if and only if s0 |=ΦrC P (x).
Proof. Observe that Bsx = Bsx entails that for any variable xi ,
s(xi ) = nil iff s0 (xi ) = nil . The result follows by verifying
that satisfaction of predicates T, F, N by a stack s depends only
on whether their arguments evaluate to nil or not.
We now present the reduction from SCVP to PREDSAT.
Definition 4.14. Let SC = (L, R) be an instance of SCVP. Let
ΦrL , ΦrR be as in Definition 4.11 (we assume names PL , PR for
the corresponding P predicate). Define rules (U1), (U2), (U3), and
inductive rule sets X and ΦSC as follows.
num1 (x) ∗ T (v) ⇒ U (x, v)
num2 (x) ∗ F (v) ⇒ U (x, v)
PL (x, l) ∗ PR (x, r) ∗
⇒ U (x, v)
U (l, w) ∗ U (r, z) ∗ N (w, z, v)
X =def {(U1), (U2), (U3)} ∪ ΦrL ∪ ΦrR
U (x, v) ∗ T (v) ⇒ Q> (x)
U (x, v) ∗ F (v) ⇒ Q⊥ (x) ∪ X
ΦSC =def
num2n (x) ∗ Q> (x) ⇒ R
Thus SC is mapped to the PREDSAT instance (ΦSC , R).
Theorem 4.15. PREDSAT is EXPTIME-complete.
Proof. By Lemma 4.6, PREDSAT is in EXPTIME. Hardness follows by exhibiting a reduction from SCVP (which is EXPTIMEcomplete [23]) via Definition 4.14. This follows from the stronger
fact that for all i ∈ {1, . . . , 2n } and for any boolean b, the output of
gate i is b iff there is a stack s such that s |=ΦSC numi (x) ∗ Qb (x).
We use induction on i.
(⇒) Assume i = 1. By assumption, gate 1 is the constant >
thus its output is also b = >. Let B =def iB . Define the
stack sˆ =def sxB [v 7→ sval(>)]. By applying (U1) we obtain that
sˆ |=X U (x, v) and as sˆ |=ΦSC T (v) by construction, it must be
that sˆ |=ΦSC numi (x) ∗ Q> (x). The case where i = 2 is similar.
Suppose i > 2 and that the output of gate i is b. Gate i is
a NAND gate thus there are l, r < i such that L(iB , lB ) = >,
R(iB , rB ) = >, the values of gates l, r are c, d and c ↑ d = b.
By the inductive hypothesis we obtain two stacks sl , sr with
sl |=ΦSC numl (l) ∗ Qc (l) and sr |=ΦSC numr (r) ∗ Qd (r). In
particular, regardless of the values c, d take, sl |=ΦSC U (l, w)
and sr |=ΦSC U (r, z) (we set sl (w) = sl (v) and sr (z) = sr (v)
w.l.o.g.). By Lemma 4.12, there exist two stacks s0l , s0r such that
Solved in
Solved in
≤ 1s
≤ 30 s
≤ 1s
≤ 30 s
≤ 1s
≤ 30 s
≤ 1s
≤ 30 s
s0l |=ΦrL PL (x, l) ∗ numi (x) ∗ numl (l)
s0r |=ΦrR PR (x, r) ∗ numi (x) ∗ numr (r)
where Lemma 4.13 lets us assume s0l (x) = s0r (x), s0l (l) = sl (l)
and s0r (r) = sr (r). Thus there exists sˆ such that sˆ(w) = sl (w),
sˆ(z) = sr (z) and
sˆ |=ΦSC PL (x, l) ∗ PR (x, r) ∗ U (l, w) ∗ U (r, z) ∗ N (w, z, v)
and by (U3), sˆ |= U (x, v). A case analysis on the definition of N
allows us to conclude that sˆ |=ΦSC numi (x) ∗ Qb (x).
(⇐) Let s be a stack such that s |=ΦSC numi (x) ∗ Qb (x). Thus
s |=ΦSC U (x, v). If i ∈ {1, 2} then gate i is a constant and the
result follows from (U1), (U2).
If i > 2 then gate i is a NAND gate and rule (U3) applies:
s |=ΦSC PL (x, l) ∗ PR (x, r) ∗ U (l, w) ∗ U (r, z) ∗ N (w, z, v).
Let l, r ∈ {1, . . . , 2n } be such that lB = Bsl and rB = Bsr .
By Lemma 4.12 we can conclude that L(iB , lB ) = > and
R(iB , rB ) = >, meaning that l, r are the inputs of gate i. As such,
there exist booleans c, d such that s |=ΦSC Qc (l) ∗ Qd (r) and by
the inductive hypothesis we obtain that the outputs of gates l, r are
c, d respectively. It remains to show that c ↑ d = b which follows
by a case analysis on the definition of N .
The worst case can be exhibited easily through counting.
Proposition 4.16. There is a family of predicates Φn of size O(n)
such that the algorithm in Defn. 3.5 runs in Ω(2n ) time and space.
Proof. It should be clear that circuits can be rendered as predicates
in linear space. In particular, the successor relation over two n-bit
numbers can be encoded as a predicate succ(x, y), for n-variable
tuples x, y. Consider the set of predicates Φn .
num1 (y) ⇒ Q(y)
succ(x, y) ∗ Q(x) ⇒ Q(y)
Φn =def
num(2n ) (x) ∗ Q(x) ⇒ P
Given (Φn , P ), the algorithm starts with the base case for Q and
counts from 1 to 2n , creating a base pair encoding each number
in between. This will happen irrespective of which strategy is used
to select rules or base pairs for instantiation. Clearly, the algorithm
will take Ω(2n ) time.
Implementation and experiments
We implemented our decision procedure as a straightforward rendition of Definition 3.5 (in about 1000 lines of OCaml code) in the
theorem proving framework C YCLIST [11], which provides support
for logics with inductive predicates. Here we report on the performance of this algorithm on various data sets.
The code, the tool and test files are available online at [1].
Benchmarks on handwritten predicates
First, we collected a set of 17 standard predicates (defined by 36
rules) from the separation logic literature, defining data structures
such as singly- or doubly-linked list segments, possibly-cyclic lists,
trees and tree segments. Our decision procedure takes just 4 ms to
prove satisfiability of this set of definitions. This is rather to be
expected, since all predicates in this test set are easily seen to be
In order to exhibit the worst-case behaviour of our algorithm,
we also tested our tool on the family Φn of predicates given in
Solved in
Solved in
Figure 2. Synthetic benchmark performance of our algorithm.
Proposition 4.16, using a binary adder implementation of the successor relation. Runtimes, as expected, are exponential in n: cases
with n ≤ 4 are solved in a fraction of a second, the case n = 5 is
solved in about half a minute, n = 6 takes almost 16 minutes, and
n = 7 times out after 40 minutes.
Benchmarks on automatically abduced predicates
Next, we tested our algorithm on a large collection of predicates
generated automatically by the C ABER tool, which attempts to abduce inductively defined safety and/or termination preconditions
in our logical fragment for pointer programs [10]. We recorded
all candidate preconditions generated by C ABER over 29 test programs, including back-tracking attempts. This produced 45 945
syntactically unique inductive rule sets, defining from 1–37 predicates with up to 11 parameters and two recursive invocations each.
The majority of these definition sets (94%) had no more than 20
predicates; sets with 19 predicates were the most numerous (15%).
We found that our algorithm terminates in less than 50 ms on all
tests in this suite, and most definition sets (83%) were satisfiable.
We believe this is probably due to the relatively simple recursive
structure of the predicates produced by C ABER.
Synthetic benchmarks
To evaluate the scalability of our procedure, we generated test cases
drawn from a random distribution with parameters:
• Vars: the number of variables in Var;
• Rules = Preds × Cases: the number of predicates and induc-
tive rules for each predicate;
• Arity: the arity of all predicates;
• Eqs, Neqs, Points, and Recs: the average numbers of equal-
ities, disequalities, points-to literals, and recursive predicate
calls, respectively, in each rule.
Each instance consists of Rules = Preds × Cases independently
generated inductive definitions, with all predicates of the same
Arity. On the rule body, the number for each kind of literal is
drawn from a Poisson distribution where the parameter λ is set
to, respectively, Eqs, Neqs, Points, and Recs. This yields a mix
of short and long rule bodies with a specified average length.
Furthermore, to have on average one base case for the rule system,
with probability p = 1/Rules all the recursive calls in the body of
a rule are discarded.
Figure 2 reports our algorithm’s performance on instances
drawn from this distribution. Each row collects the results of 100
instances randomly generated with various parameter values; the
second and third columns report the number of tests solved in under 1 and 30 seconds, while the last one shows the number of solved
instances found to be satisfiable. Each of the four sub-tables shows
how the statistics vary as one parameter changes while all the rest
remain fixed. Although most cases are fairly easy to solve, a few
hard instances consistently show up on all parameter settings. For
example, with parameters Vars = 50, Rules = 3×3, Arity = 3,
and all Eqs, Neqs, Points, and Recs set to 2 (that is the bold line
repeated on all four tables), we find that two very hard instances
remain unsolved after the 30 s timeout.
Related work
Here we survey the main categories of work related to the contribution of this paper.
Satisfiability in separation logic with inductive predicates
Recently, interesting progress has been made on satisfiability and
entailment problems in various fragments of separation logic with
user-defined inductive predicates as considered here (see Section 2). In particular, Iosif et al. [20] propose a sub-fragment of
our setting, imposing significant syntactic restrictions on the inductive definitions to enforce bounded treewidth for all models
of a predicate. Here the treewidth of a heap h is determined by
a corresponding graph structure. These restrictions allow them to
use a reduction to monadic second-order logic (MSO) for their
proof that both satisfiability and entailment are decidable in their
fragment of separation logic with inductive definitions. However,
these syntactic restrictions disallow many natural inductive predicates; in particular, predicates describing structures with dangling
data pointers, such as list or tree segments with extra arbitrary data
fields, or structures where potentially no memory is allocated.1
In this paper, we only consider satisfiability, not entailment
(which is undecidable [2]), but we are not restricted to boundedtreewidth predicates. In addition, we provide a direct decision procedure for satisfiability (as opposed to a reduction proof of decidability) and contribute an analysis of the complexity of satisfiability
checking for our fragment.
Considerable research effort has also been expended on the
symbolic heap fragment of separation logic with a list segment
predicate only. Berdine et al. [3] provided the first decidability result for satisfiability and entailment in this fragment, with a polynomial time procedure given later by Cook et al. [15]. Piskac et al.
recently presented another decision procedure, which also works
for structures such as sorted list segments and doubly linked lists,
based on translating entailments to an intermediate logic that can be
handled by an SMT solver [24]. (In very recent work [25], Piskac et
al. extended their approach to support also tree-shaped data structures.) Finally, Navarro P´erez and Rybalchenko [22] provided an
SMT encoding for satisfiability and entailment in another extension of the fragment allowing pure formulas from arbitrary SMT
theories, rather than simple (dis)equalities.
Satisfiability in other logics with inductive definitions
The evaluation of Datalog queries was shown to be decidable by
Shmueli [27] and EXPTIME-complete (see e.g. [16, Theorem
4.5]). However, the interaction between spatial conjunction (∗),
allocation (7→) and recursion makes a direct translation to Datalog impractical because it imposes a global constraint; at the same
time we found that reducing from Datalog for our lower bounds is
possible but no simpler than our circuit-based method.
More recently, Hoder et al. [19] describe a Datalog-based engine, µZ, for the fixed point computation of inductive definitions.
In contrast to our approach, they compute concrete fixed points
based on explicit underlying ground facts (i.e., they focus on model
checking, as opposed to satisfiability checking).
More generally, there has been some interest from the program
analysis and verification community in the satisfiability of Horn
and Horn-like clauses. Bjørner et al. [7] advocate such clauses as
an interchange format for software model checking tools, while
Grebenshchikov et al. [18] describe an abstraction-based procedure
for finding models and checking satisfiability of Horn-like clauses
in the context of program analysis.
Separation logic tools with general inductive predicates
Several analysis tools based on separation logic allow the user
to provide their own inductive definitions for spatial predicates.
We have already mentioned in our evaluation (Section 5) the theorem prover C YCLIST [11], which treats separation logic with
user-defined inductive predicates, and the related abductive prover
C ABER [10], which automatically infers inductive predicate definitions as safety/termination preconditions for while programs. Another such tool is T HOR [21], which proves memory safety and generates sound arithmetic abstractions of heap programs (w.r.t. safety
and termination). In T HOR the specification and predicate definitions must be entered manually. User-defined inductive predicates
are also employed in H IP /S LEEK [14], a combined theorem prover
and verification system for a C-like language. Finally, the shape
analysis in [13] infers inductive definitions based on “structural invariant checkers” provided by the user.
Conclusion and future work
The decidability status of satisfiability in the symbolic heap fragment of separation logic with general inductive predicates has stood
open for some time. Following the recent achievement of a partial
positive answer to this question in [20], here we resolve the general case affirmatively. Our decidability proof has the advantage
of being constructive: we give a decision procedure for checking
the satisfiability of inductively defined predicates and of individual
rules defining those predicates.
We show that our satisfiability problem is EXPTIME-complete
in the general case, and that it is still NP-complete when the inductive predicates are restricted to at most k ≥ 3 arguments. Despite
these high complexities, our experiments indicate that, for predicate definitions arising in practice, our prototype implementation is
typically able to solve the decision problem in a matter of milliseconds. This opens up a number of interesting potential applications
in the automatic verification of programs based on our fragment of
separation logic (which could be used, e.g., to consider heap-based
programs employing arbitrary data structures, rather than just lists):
The consistency of inductive predicates defined by first-order Horn
clauses has been widely studied in different contexts. One wellknown application of such clausal definitions is Datalog, a rulebased query language for relational databases with ties to logic
programming. Datalog rules roughly correspond to our inductive
rules where the spatial component of rule bodies is always emp.
• First, automatic verification or shape analysis based on sepa-
1 For
• Second, although our decision procedure for satisfiability can-
further discussion of the limitations of this fragment of separation
logic we refer to [20].
ration logic (cf. [4, 5, 12]) typically employs symbolic states
based on disjunctions of symbolic heaps. Our algorithm could
thus be used to quickly eliminate unsatisfiable disjuncts from
symbolic states; it seems plausible that this might yield nontrivial reductions in the time and space costs of such analyses.
not be used to decide the validity of entailments (which is im-
possible in general), it does nevertheless provide a quick partial
method for proving or disproving such entailments. On the one
hand, any entailment F ` G with an unsatisfiable antecedent F
is trivially valid. On the other hand, F ` G must be invalid if
the base pairs for F and G imply the existence of a model satisfying F but not G. To illustrate this point, consider the usual
“list segment” predicate of separation logic, defined by
x = y : emp ⇒ ls(x, y)
x 7→ z ∗ ls(z, y) ⇒ ls(x, y)
Now, consider the (invalid) entailment ls(x, y) ` ls(y, x). Computing the base pairs for both sides yields
(base ls)(x, y) = {(∅, {x = y}), ({x}, ∅)}
(base ls)(y, x) = {(∅, {y = x}), ({y}, ∅)}
The second base pair for ls(x, y) implies the existence of a
model for ls(x, y) in which x is allocated and y is not, and
hence x 6= y in this model. However, the base pairs for ls(y, x)
tell us that y must be allocated in any model of ls(y, x) where
x 6= y. That is, there is a model of ls(x, y) that is not a model of
ls(y, x). We believe that this sort of reasoning should be quite
straightforward to automate.
• Third, our satisfiability procedure can be used to check the san-
ity of, and/or remove unsatisfiable clauses from, the definitions
of inductive predicates either written by programmers, or inferred automatically (cf. [10]).
• Finally, beyond deciding the satisfiability of inductive predi-
cates, the set of base pairs computed by our procedure also provides a useful partition of the space to search for models. While
Gherghina et al. [17] have shown that a manual case analysis
of inductive definitions can guide proof techniques and improve
their efficiency; our decision procedure could automate the case
analysis required to do so.
Taking a different direction for future work, one could investigate whether our techniques for deciding satisfiability extend to
more general variants of separation logic, where general inductive predicates are still allowed, but the general format of formulas
is less restricted. Possible candidates for such extensions include:
higher-order separation logic [6]; the fragment in which formulas
may contain pure assertions beyond (dis)equalities [22]; and separation logic with fractional permissions [8].
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