Verification conditions

Software verification aims at making software safer. In its typical use case, it is a tool that takes as input the source code of a program with specifications as annotations and attempt to prove — or disprove — their validity.

One of the core modules of Stainless is a verifier for the subset of Scala described in the sections Pure Scala and Imperative. In this section, we describe the specification language that can be used to declare properties of programs, as well as the safety properties automatically checked by Stainless. We also discuss how Stainless can be used to prove mathematical theorems.

Given an input program, Stainless generates individual verification conditions corresponding to different properties of the program. A program is correct if all of the generated verification conditions are valid. The validity of some conditions depends on the validity of other conditions — typically a postcondition is valid assuming the precondition is valid.

For each function, Stainless attempts to verify its contract, if there is one. A contract is the combination of a precondition and a postcondition. A function meets its contract if, for any input parameter that passes the precondition, the postcondition holds after executing the function. Preconditions and postconditions are annotations given by the user — they are the specifications and hence cannot be inferred by a tool and must be provided.

In addition to user-provided contracts, Stainless will also generate a few safety verification conditions of its own. It will check that all of the array accesses are within proper bounds, and that pattern matching always covers all possible cases, even given complex precondition. The latter is different from the Scala compiler checks on pattern matching exhaustiveness, as Stainless considers information provided by (explicit or implicit) preconditions to the match expression.


One core concept in verification is to check the contract of functions. The most important part of a contract is the postcondition. The postcondition specifies the behavior of the function. It takes into account the precondition and only asserts the property of the output when the input satisfies the precondition.

Formally, we consider a function with a single parameter (one can generalize the following for any number of parameters):

def f(x: A): B = {
} ensuring(r => post)

where, \(\mbox{prec}(x)\) is a Boolean expression with free variables contained in \(\{ x \}\), \(\mbox{body}(x)\) is an expression with free variables contained in \(\{ x \}\) and \(\mbox{post}(x, r)\) is a Boolean expression with free variables contained in \(\{ x, r \}\). The types of \(x\) and \(r\) are respectively A and B. We write \(\mbox{expr}(a)\) to mean the substitution in \(\mbox{expr}\) of its free variable by \(a\).

Stainless attempts to prove the following theorem:

\[\forall x. \mbox{prec}(x) \implies \mbox{post}(x, \mbox{body}(x))\]

If Stainless is able to prove the above theorem, it returns valid for the function f. This gives you a guarantee that the function f is correct with respect to its contract.

However, if the theorem is not valid, Stainless will return a counterexample to the theorem. The negation of the theorem is:

\[\exists x. \mbox{prec}(x) \land \neg \mbox{post}(x, \mbox{body}(x))\]

and to prove the validity of the negation, Stainless finds a witness \(x\) — a counterexample — such that the precondition is verified and the postcondition is not.

The general problem of verification is undecidable for a Turing-complete language, and the Stainless language is Turing-complete. So Stainless has to be incomplete in some sense. Generally, Stainless will eventually find a counterexample if one exists. However, in practice, some program structures require too many unrollings and Stainless is likely to timeout (or being out of memory) before finding the counterexample. When the postcondition is valid, it could also happen that Stainless keeps unrolling the program forever, without being able to prove the correctness. We discuss the exact conditions for this in the chapter on Stainless’s algorithms.


Preconditions are used as part of the contract of functions. They are a way to restrict the input to only relevant inputs, without having to implement guards and error handling in the functions themselves.

Preconditions are contracts that the call sites should respect, and thus are not checked as part of verifying the function. Given the following definition:

def f(x: A): B = {

For each call site in the whole program (including in f itself):


where the expression \(\mbox{e}(x)\) is an expression of type A with free variables among \(\{ x \}\). Let us define the path condition on \(x\) at that program point as \(\mbox{pc}(x)\). The path condition is a formula that summarizes the facts known about \(x\) at that call site. A typical example is when the call site is inside an if expression:

if(x > 0)

The path condition on \(x\) would include the fact that \(x > 0\). This path condition is then used as a precondition of proving the validity of the call to \(\mbox{f}\). Formally, for each such call site, Stainless will attempt to prove the following theorem:

\[\forall x. \mbox{pc}(x) \implies \mbox{prec}(\mbox{e}(x))\]

Stainless will generate one such theorem for each static call site of a function with a precondition.


Stainless only assumes an open program model, where any function could be called from outside of the given program. In particular, Stainless will not derive a precondition to a function based on known information in the context of the calls, such as knowing that the function is always given positive parameters. Any information needed to prove the postcondition will have to be provided as part of the precondition of a function.

Sharing bindings between specifications and function body

The example ValEnsuring shows that Stainless supports multiple require’s (in functions, but not for ADT invariants), and shows how to share a val binding between precondition, postcondition, and function body.

Loop invariants

Stainless supports annotations for loop invariants in Imperative. To simplify the presentation we will assume a single variable \(x\) is in scope, but the definitions generalize to any number of variables. Given the following program:

(while(cond) {
}) invariant(inv)

where the Boolean expression \(\mbox{cond}(x)\) is over the free variable \(x\) and the expression \(\mbox{body}(x, x')\) relates the value of \(x\) when entering the loop to its updated value \(x'\) on loop exit. The expression \(\mbox{inv}(x)\) is a Boolean formula over the free variable \(x\).

A loop invariant must hold:
  1. when the program enters the loop initially

  2. after each completion of the body

  3. right after exiting the loop (when the condition turns false)

Stainless will prove the points (1) and (2) above. Together, and by induction, they imply that point (3) holds as well.

Stainless also supports noReturnInvariant (see ReturnInWhile3) to describe loop invariants that are allowed to be broken after a return (can be combined with invariant).

Decrease annotation in loops

One can also specify that the value of a given expression of numerical type decreases at each loop iteration by adding a decreases measure within the loop body:

while(cond) {

Stainless will then emit a verification condition that checks whether the expression is strictly positive and decreases at each iteration.

Array access safety

Stainless generates verification conditions for the safety of array accesses. For each array variable, Stainless carries along a symbolic information on its length. This information is used to prove that each expression used as an index in the array is strictly smaller than that length. The expression is also checked to be positive.

ADT invariants

Stainless lets the user write ADT invariants with the require keyword. Internally, such invariants are extracted as methods (named inv). Whenever, an ADT is constructed, and to make sure that the invariant is true, a verification condition is generated with a call to the corresponding inv method.

The Stainless annotation @inlineInvariant used on an ADT or one of its ancestors can be used to inline the call to inv in the verification condition, as shown in the following example. This annotation is only supported when --type-checker=true (which is the case by default).

import stainless.annotation._

object InlineInvariant {
  sealed abstract class A

  case class B(x: BigInt) extends A {
    require(x >= 50)

  case class C(x: BigInt) extends A {
    require(x >= 50)

  def f(): A = {
    B(100) // VC: inv(B(100))
    c(100) // VC: 100 >= 50 (call to `inv` was inlined)

Pattern matching exhaustiveness

Stainless verifies that pattern matching is exhaustive. The regular Scala compiler only considers the types of expression involved in pattern matching, but Stainless will consider information such as precondition to formally prove the exhaustiveness of pattern matching.

As an example, the following code should issue a warning with Scala:

abstract class List
case class Cons(head: Int, tail: List) extends List
case object Nil extends List

def getHead(l: List): Int = {
  l match {
    case Cons(x, _) => x

But Stainless will prove that the pattern matching is actually exhaustive, relying on the given precondition.