# HoareHoare Logic

(* \$Date: 2011-06-22 14:56:13 -0400 (Wed, 22 Jun 2011) \$ *)

(* Alexandre Pilkiewicz suggests the following alternate type for the
decorated WHILE construct:
| DCWhile : bexp -> Assertion -> dcom -> Assertion -> dcom
This leads to a simpler rule in the VC generator, which is much
easier to explain:
| DCWhile b P' c  => ((fun st => post c st /\ bassn b st) ~~> P')
/\ (P ~~> post c)  (* post c is the loop invariant *)
/\ verification_conditions P' c
His full development (based on an old version of our formalized
decorated programs, unfortunately), can be found in the file
/underconstruction/PilkiewiczFormalizedDecorated.v *)

Require Export ImpList.

We've begun applying the mathematical tools developed in the first part of the course to studying the theory of a small programming language, Imp.
• We defined a type of abstract syntax trees for Imp, together with an evaluation relation (a partial function on states) that specifies the operational semantics of programs.
The language we defined, though small, captures some of the key features of full-blown languages like C, C++, and Java, including the fundamental notion of mutable state and some common control structures.
• We proved a number of metatheoretic properties — "meta" in the sense that they are properties of the language as a whole, rather than properties of particular programs in the language. These included:
• determinacy of evaluation
• equivalence of some different ways of writing down the definition
• guaranteed termination of certain classes of programs
• correctness (in the sense of preserving meaning) of a number of useful program transformations
• behavioral equivalence of programs (in the optional chapter Equiv.v).
If we stopped here, we would still have something useful: a set of tools for defining and discussing programming languages and language features that are mathematically precise, flexible, and easy to work with, applied to a set of key properties.
All of these properties are things that language designers, compiler writers, and users might care about knowing. Indeed, many of them are so fundamental to our understanding of the programming languages we deal with that we might not consciously recognize them as "theorems." But properties that seem intuitively obvious can sometimes be quite subtle — or, in some cases, actually even wrong!
We'll return to this theme later in the course when we discuss types and type soundness.
• We saw a couple of examples of program verification — using the precise definition of Imp to prove formally that certain particular programs (e.g., factorial and slow subtraction) satisfied particular specifications of their behavior.
In this chapter, we'll take this last idea further. We'll develop a reasoning system called Floyd-Hoare Logic — commonly, if somewhat unfairly, shortened to just Hoare Logic — in which each of the syntactic constructs of Imp is equipped with a single, generic "proof rule" that can be used to reason about programs involving this construct.
Hoare Logic originates in the 1960s, and it continues to be the subject of intensive research right up to the present day. It lies at the core of a huge variety of tools that are now being used to specify and verify real software systems.

# Hoare Logic

Hoare Logic offers two important things: a natural way of writing down specifications of programs, and a compositional proof technique for proving that these specifications are met — where by "compositional" we mean that the structure of proofs directly mirrors the structure of the programs that they are about.

## Assertions

If we're going to talk about specifications of programs, the first thing we'll want is a way of making assertions about properties that hold at particular points in time — i.e., properties that may or may not be true of a given state of the memory.

Definition Assertion := state Prop.

#### Exercise: 1 star (assertions)

Paraphrase the following assertions in English.
fun st =>  asnat (st X) = 3
fun st =>  asnat (st X) = x
fun st =>  asnat (st X) <= asnat (st Y)
fun st =>  asnat (st X) = 3  asnat (st X) <= asnat (st Y)
fun st =>  (asnat (st Z)) * (asnat (st Z)) <= x
~ (((S (asnat (st Z))) * (S (asnat (st Z)))) <= x)
fun st =>  True
fun st =>  False
This way of writing assertions is formally correct — it precisely captures what we mean, and it is exactly what we will use in Coq proofs — but it is a bit heavy to look at, for several reasons. First, every single assertion that we ever write is going to begin with fun st => ; (2) this state st is the only one that we ever use to look up variables (we never need to talk about two different states at the same time); and (3) all the variable lookups in assertions are cluttered with asnat or aslist coercions. When we are writing down assertions informally, we can make some simplifications: drop the initial fun st =>, write just X instead of st X, and elide asnat and aslist. Informally, instead of writing
fun st =>  (asnat (st Z)) * (asnat (st Z)) <= x
~ ((S (asnat (st Z))) * (S (asnat (st Z))) <= x)
we'll write just
Z * Z <= x
~((S Z) * (S Z) <= x).

## Hoare Triples

Next, we need a way of specifying — making claims about — the behavior of commands.
Since we've defined assertions as a way of making claims about the properties of states, and since the behavior of a command is to transform one state to another, a claim about a command takes the following form:
• "If c is started in a state satisfying assertion P, and if c eventually terminates, then the final state is guaranteed to satisfy the assertion Q."
Such a claim is called a Hoare Triple. The property P is called the precondition of c; Q is the postcondition of c.
(Traditionally, Hoare triples are written {P} c {Q}, but single braces are already used for other things in Coq.)

Definition hoare_triple (P:Assertion) (c:com) (Q:Assertion) : Prop :=
st st',
c / st st'
P st
Q st'.

Since we'll be working a lot with Hoare triples, it's useful to have a compact notation:
{{P}}  c  {{Q}}.

Notation "{{ P }} c" := (hoare_triple P c (fun st => True)) (at level 90)
: hoare_spec_scope.
Notation "{{ P }} c {{ Q }}" := (hoare_triple P c Q) (at level 90, c at next level)
: hoare_spec_scope.
Open Scope hoare_spec_scope.

(The hoare_spec_scope annotation here tells Coq that this notation is not global but is intended to be used in particular contexts. The Open Scope tells Coq that this file is one such context. The first notation — with missing postcondition — will not actually be used here; it's just a placeholder for a notation that we'll want to define later, when we discuss decorated programs.)

#### Exercise: 1 star (triples)

Paraphrase the following Hoare triples in English.
{{True}} c {{X = 5}}

{{X = x}} c {{X = x + 5)}}

{{X <= Y}} c {{Y <= X}}

{{True}} c {{False}}

{{X = x}}
c
{{Y = real_fact x}}.

{{True}}
c
{{(Z * Z) <= x  ~ (((S Z) * (S Z)) <= x)}}

#### Exercise: 1 star (valid_triples)

Which of the following Hoare triples are valid — i.e., the claimed relation between P, c, and Q is true?
{{True}} X ::= 5 {{X = 5}}

{{X = 2}} X ::= X + 1 {{X = 3}}

{{True}} X ::= 5; Y ::= 0 {{X = 5}}

{{X = 2  X = 3}} X ::= 5 {{X = 0}}

{{True}} SKIP {{False}}

{{False}} SKIP {{True}}

{{True}} WHILE True DO SKIP END {{False}}

{{X = 0}}
WHILE X == 0 DO X ::= X + 1 END
{{X = 1}}

{{X = 1}}
WHILE X <> 0 DO X ::= X + 1 END
{{X = 100}}
(Note that we're using informal mathematical notations for expressions inside of commands, for readability. We'll continue doing so throughout the chapter.)
To get us warmed up, here are two simple facts about Hoare triples.

Theorem hoare_post_true : (P Q : Assertion) c,
( st, Q st)
{{P}} c {{Q}}.
Proof.
intros P Q c H. unfold hoare_triple.
intros st st' Heval HP.
apply H. Qed.

Theorem hoare_pre_false : (P Q : Assertion) c,
( st, ~(P st))
{{P}} c {{Q}}.
Proof.
intros P Q c H. unfold hoare_triple.
intros st st' Heval HP.
unfold not in H. apply H in HP.
inversion HP. Qed.

## Weakest Preconditions

Some Hoare triples are more interesting than others. For example,
{{ False }}  X ::= Y + 1  {{ X <= 5 }}
is not very interesting: it is perfectly valid, but it tells us nothing useful. Since the precondition isn't satisfied by any state, it doesn't describe any situations where we can use the command X ::= Y + 1 to achieve the postcondition X <= 5.
By contrast,
{{ Y <= 4  Z = 0 }}  X ::= Y + 1 {{ X <= 5 }}
is useful: it tells us that, if we can somehow create a situation in which we know that Y <= 4 Z = 0, then running this command will produce a state satisfying the postcondition. However, this triple is still not as useful as it could be, because the Z = 0 clause in the precondition actually has nothing to do with the postcondition X <= 5. The most useful triple (with the same command and postcondition) is this one:
{{ Y <= 4 }}  X ::= Y + 1  {{ X <= 5 }}
In other words, Y <= 4 is the weakest valid precondition of the command X ::= Y + 1 for the postcondition X <= 5.
In general, we say that "P is the weakest precondition of c for Q" if
• {{P}} c {{Q}}, and
• whenever P' is an assertion such that {{P'}} c {{Q}}, we have P' st implies P st for all states st.
That is, P is the weakest precondition of c for Q if (a) P is a precondition for Q and c, and (b) P is the weakest (easiest to satisfy) assertion that guarantees Q after c.

#### Exercise: 1 star (wp)

What are the weakest preconditions of the following commands for the following postconditions?
{{ ? }}  SKIP  {{ X = 5 }}

{{ ? }}  X ::= Y + Z {{ X = 5 }}

{{ ? }}  X ::= Y  {{ X = Y }}

{{ ? }}
IFB X == 0 THEN Y ::= Z + 1 ELSE Y ::= W + 2 FI
{{ Y = 5 }}

{{ ? }}
X ::= 5
{{ X = 0 }}

{{ ? }}
WHILE True DO X ::= 0 END
{{ X = 0 }}

## Proof Rules

The goal of Hoare logic is to provide a compositional method for proving the validity of Hoare triples. That is, the structure of a program's correctness proof should mirror the structure of the program itself. To this end, in the sections below, we'll introduce one rule for reasoning about each of the different syntactic forms of commands in Imp — one for assignment, one for sequencing, one for conditionals, etc. — plus a couple of "structural" rules that are useful for gluing things together.

### Assignment

The rule for assignment is the most fundamental of the Hoare logic proof rules. Here's how it works.
Consider this (valid) Hoare triple:
{{ Y = 1 }}  X ::= Y  {{ X = 1 }}
In English: if we start out in a state where the value of Y is 1 and we assign Y to X, then we'll finish in a state where X is 1. That is, the property of being equal to 1 gets transferred from Y to X.
Similarly, in
{{ Y + Z = 1 }}  X ::= Y + Z  {{ X = 1 }}
the same property (being equal to one) gets transferred to X from the expression Y + Z on the right-hand side of the assignment.
More generally, if a is any arithmetic expression, then
{{ a = 1 }}  X ::= a {{ X = 1 }}
is a valid Hoare triple.
Even more generally, a is any arithmetic expression and Q is any property of numbers, then
{{ Q(a) }}  X ::= a {{ Q(X) }}
is a valid Hoare triple.
Rephrasing this a bit gives us the general Hoare rule for assignment:
{{ Q where a is substituted for X }}  X ::= a  {{ Q }}
For example, these are valid applications of the assignment rule:
{{ X + 1 <= 5 }}  X ::= X + 1  {{ X <= 5 }}

{{ 3 = 3 }}  X ::= 3  {{ X = 3 }}

{{ 0 <= 3  3 <= 5 }}  X ::= 3  {{ 0 <= X  X <= 5 }}
We could try to formalize the assignment rule directly in Coq by treating Q as a family of assertions indexed by arithmetic expressions — something like this:
Theorem hoare_asgn_firsttry :
(Q : aexp  AssertionV a,
{{fun st => Q a st}} (V ::= a) {{fun st => Q (AId Vst}}.
But this formulation is not very nice, for two reasons. First, it is not quite true! (As a counterexample, consider a Q that inspects the syntax of its argument, such as
Definition Q (a:aexp) : Prop :=
match a with
| AID (Id 0) => fun st => False
| _          => fun st => True
end.
together with any V = Id 0 because a precondition that reduces to True leads to a postcondition that is False.) And second, even if we could prove something similar to this, it would be awkward to use.
A much smoother way of formalizing the rule arises from the following observation:
• "Q where a is substituted for X" holds in a state st iff Q holds in the state update st X (aeval st a).
That is, asserting that a substituted variant of Q holds in some state is the same as asserting that Q itself holds in a substituted variant of the state.
Here is the definition of substitution in a state:

Definition assn_sub V a Q : Assertion :=
fun (st : state) =>
Q (update st V (aeval st a)).

This gives us the formal proof rule for assignment:
 (hoare_asgn) {{assn_sub V a Q}} V::=a {{Q}}

Theorem hoare_asgn : Q V a,
{{assn_sub V a Q}} (V ::= a) {{Q}}.
Proof.
unfold hoare_triple.
intros Q V a st st' HE HQ.
inversion HE. subst.
unfold assn_sub in HQ. assumption. Qed.

Here's a first formal proof using this rule.

Example assn_sub_example :
{{fun st => 3 = 3}}
(X ::= (ANum 3))
{{fun st => asnat (st X) = 3}}.
Proof.
assert ((fun st => 3 = 3) =
(assn_sub X (ANum 3) (fun st => asnat (st X) = 3))).
Case "Proof of assertion".
unfold assn_sub. reflexivity.
rewrite H. apply hoare_asgn. Qed.

This proof is a little clunky because the hoare_asgn rule doesn't literally apply to the initial goal: it only works with triples whose precondition has precisely the form assn_sub Q V a for some Q, V, and a. So we have to start with asserting a little lemma to get the goal into this form.
Doing this kind of fiddling with the goal state every time we want to use hoare_asgn would get tiresome pretty quickly. Fortunately, there are easier alternatives. One simple one is to state a slightly more general theorem that introduces an explicit equality hypothesis:

Theorem hoare_asgn_eq : Q Q' V a,
Q' = assn_sub V a Q
{{Q'}} (V ::= a) {{Q}}.
Proof.
intros Q Q' V a H. rewrite H. apply hoare_asgn. Qed.

With this version of hoare_asgn, we can do the proof much more smoothly.

Example assn_sub_example' :
{{fun st => 3 = 3}}
(X ::= (ANum 3))
{{fun st => asnat (st X) = 3}}.
Proof.
apply hoare_asgn_eq. reflexivity. Qed.

#### Exercise: 2 stars (hoare_asgn_examples)

Translate these informal Hoare triples...
{{ X + 1 <= 5 }}  X ::= X + 1  {{ X <= 5 }}
{{ 0 <= 3  3 <= 5 }}  X ::= 3  {{ 0 <= X  X <= 5 }}
...into formal statements and use hoare_asgn_eq to prove them.

(* FILL IN HERE *)

#### Exercise: 3 stars (hoarestate2)

The assignment rule looks backward to almost everyone the first time they see it. If it still seems backward to you, it may help to think a little about alternative "forward" rules. Here is a seemingly natural one:
{{ True }} X ::= a {{ X = a }}
Explain what is wrong with this rule.
(* FILL IN HERE *)

#### Exercise: 3 stars, optional (hoare_asgn_weakest)

Show that the precondition in the rule hoare_asgn is in fact the weakest precondition.

Theorem hoare_asgn_weakest : P V a Q,
{{P}} (V ::= a) {{Q}}
st, P st assn_sub V a Q st.
Proof.
(* FILL IN HERE *) Admitted.

### Consequence

The discussion above about the awkwardness of applying the assignment rule illustrates a more general point: sometimes the preconditions and postconditions we get from the Hoare rules won't quite be the ones we want — they may (as in the above example) be logically equivalent but have a different syntactic form that fails to unify with the goal we are trying to prove, or they actually may be logically weaker (for preconditions) or stronger (for postconditions) than what we need.
For instance, while
{{3 = 3}} X ::= 3 {{X = 3}},
follows directly from the assignment rule, the more natural triple
{{True}} X ::= 3 {{X = 3}}.
does not. This triple is also valid, but it is not an instance of hoare_asgn (or hoare_asgn_eq) because True and 3 = 3 are not syntactically equal assertions.
In general, if we can derive {{P}} c {{Q}}, it is valid to change P to P' as long as P' is strong enough to imply P, and change Q to Q' as long as Q implies Q'.
This observation is captured by the following Rule of Consequence.
 {{P'}} c {{Q'}} P implies P' (in every state) Q' implies Q (in every state) (hoare_consequence) {{P}} c {{Q}}
For convenience, we can state two more consequence rules — one for situations where we want to just strengthen the precondition, and one for when we want to just weaken the postcondition.
 {{P'}} c {{Q}} P implies P' (in every state) (hoare_consequence_pre) {{P}} c {{Q}}
 {{P}} c {{Q'}} Q' implies Q (in every state) (hoare_consequence_post) {{P}} c {{Q}}
Here are the formal versions:

Theorem hoare_consequence : (P P' Q Q' : Assertion) c,
{{P'}} c {{Q'}}
( st, P st P' st)
( st, Q' st Q st)
{{P}} c {{Q}}.
Proof.
intros P P' Q Q' c Hht HPP' HQ'Q.
intros st st' Hc HP.
apply HQ'Q. apply (Hht st st'). assumption.
apply HPP'. assumption. Qed.

Theorem hoare_consequence_pre : (P P' Q : Assertion) c,
{{P'}} c {{Q}}
( st, P st P' st)
{{P}} c {{Q}}.
Proof.
intros P P' Q c Hhoare Himp.
apply hoare_consequence with (P' := P') (Q' := Q);
try assumption.
intros st H. apply H. Qed.

Theorem hoare_consequence_post : (P Q Q' : Assertion) c,
{{P}} c {{Q'}}
( st, Q' st Q st)
{{P}} c {{Q}}.
Proof.
intros P Q Q' c Hhoare Himp.
apply hoare_consequence with (P' := P) (Q' := Q');
try assumption.
intros st H. apply H. Qed.

For example, we might use (the "_pre" version of) the consequence rule like this:
{{ True }} =>
{{ 1 = 1 }}
X ::= 1
{{ X = 1 }}
Or, formally...

Example hoare_asgn_example1 :
{{fun st => True}} (X ::= (ANum 1)) {{fun st => asnat (st X) = 1}}.
Proof.
apply hoare_consequence_pre with (P' := (fun st => 1 = 1)).
apply hoare_asgn_eq. reflexivity.
intros st H. reflexivity. Qed.

### Digression: The eapply Tactic

This is a good moment to introduce another convenient feature of Coq. Having to write P' explicitly in the example above is a bit annoying because the very next thing we are going to do — applying the hoare_asgn rule — is going to determine exactly what it should be. We can use eapply instead of apply to tell Coq, essentially, "The missing part is going to be filled in later."

Example hoare_asgn_example1' :
{{fun st => True}}
(X ::= (ANum 1))
{{fun st => asnat (st X) = 1}}.
Proof.
eapply hoare_consequence_pre.
apply hoare_asgn_eq. reflexivity. (* or just apply hoare_asgn. *)
intros st H. reflexivity. Qed.

In general, eapply H tactic works just like apply H except that, instead of failing if unifying the goal with the conclusion of H does not determine how to instantiate all of the variables appearing in the premises of H, eapply H will replace these variables with existential variables (written ?nnn) as placeholders for expressions that will be determined (by further unification) later in the proof.
There is also an eassumption tactic that works similarly.

### Skip

Since SKIP doesn't change the state, it preserves any property P:
 (hoare_skip) {{ P }} SKIP {{ P }}

Theorem hoare_skip : P,
{{P}} SKIP {{P}}.
Proof.
intros P st st' H HP. inversion H. subst.
assumption. Qed.

### Sequencing

More interestingly, if the command c1 takes any state where P holds to a state where Q holds, and if c2 takes any state where Q holds to one where R holds, then doing c1 followed by c2 will take any state where P holds to one where R holds:
 {{ P }} c1 {{ Q }} {{ Q }} c2 {{ R }} (hoare_seq) {{ P }} c1;c2 {{ R }}

Theorem hoare_seq : P Q R c1 c2,
{{Q}} c2 {{R}}
{{P}} c1 {{Q}}
{{P}} c1;c2 {{R}}.
Proof.
intros P Q R c1 c2 H1 H2 st st' H12 Pre.
inversion H12; subst.
apply (H1 st'0 st'); try assumption.
apply (H2 st st'0); try assumption. Qed.

Note that, in the formal rule hoare_seq, the premises are given in "backwards" order (c2 before c1). This matches the natural flow of information in many of the situations where we'll use the rule.
Informally, a nice way of recording a proof using this rule is as a "decorated program" where the intermediate assertion Q is written between c1 and c2:
{{ a = n }}
X ::= a;
{{ X = n }}      <---- decoration for Q
SKIP
{{ X = n }}

Example hoare_asgn_example3 : a n,
{{fun st => aeval st a = n}}
(X ::= a; SKIP)
{{fun st => st X = n}}.
Proof.
intros a n. eapply hoare_seq.
Case "right part of seq".
apply hoare_skip.
Case "left part of seq".
eapply hoare_consequence_pre. apply hoare_asgn.
intros st H. subst. reflexivity. Qed.

#### Exercise: 2 stars (hoare_asgn_example4)

Translate this decorated program into a formal proof:
{{ True }} =>
{{ 1 = 1 }}
X ::= 1;
{{ X = 1 }} =>
{{ X = 1  2 = 2 }}
Y ::= 2
{{ X = 1  Y = 2 }}

Example hoare_asgn_example4 :
{{fun st => True}} (X ::= (ANum 1); Y ::= (ANum 2))
{{fun st => asnat (st X) = 1 asnat (st Y) = 2}}.
Proof.
(* FILL IN HERE *) Admitted.

#### Exercise: 3 stars, optional (swap_exercise)

Write an Imp program c that swaps the values of X and Y and show (in Coq) that it satisfies the following specification:
{{X <= Y}} c {{Y <= X}}

(* FILL IN HERE *)

#### Exercise: 3 stars, optional (hoarestate1)

Explain why the following proposition can't be proven:
(a : aexp) (n : nat),
{{fun st => aeval st a = n}} (X ::= (ANum 3); Y ::= a
{{fun st => asnat (st Y) = n}}.

(* FILL IN HERE *)

### Conditionals

What sort of rule do we want for reasoning about conditional commands? Certainly, if the same assertion Q holds after executing either branch, then it holds after the whole conditional. So we might be tempted to write:
 {{P}} c1 {{Q}} {{P}} c2 {{Q}} {{P}} IFB b THEN c1 ELSE c2 {{Q}}
However, this is rather weak. For example, using this rule, we cannot show that:
{{True}}
IFB X == 0
THEN Y ::= 2
ELSE Y ::= X + 1
FI
{{ X <= Y }}
since the rule tells us nothing about the state in which the assignments take place in the "then" and "else" branches.
But, actually, we can say something more precise. In the "then" branch, we know that the boolean expression b evaluates to true, and in the "else" branch, we know it evaluates to false. Making this information available in the premises of the lemma gives us more information to work with when reasoning about the behavior of c1 and c2 (i.e., the reasons why they establish the postcondtion Q).
 {{P ∧  b}} c1 {{Q}} {{P ∧ ~b}} c2 {{Q}} (hoare_if) {{P}} IFB b THEN c1 ELSE c2 FI {{Q}}
To interpret this rule formally, we need to do a little work.
Strictly speaking, the assertion we've written, P b, is the conjunction of an assertion and a boolean expression, which doesn't typecheck. To fix this, we need a way of formally "lifting" any bexp b to an assertion. We'll write bassn b for the assertion "the boolean expression b evaluates to true (in the given state)."

Definition bassn b : Assertion :=
fun st => (beval st b = true).

A couple of useful facts about bassn:

Lemma bexp_eval_true : b st,
beval st b = true (bassn b) st.
Proof.
intros b st Hbe.
unfold bassn. assumption. Qed.

Lemma bexp_eval_false : b st,
beval st b = false ~ ((bassn b) st).
Proof.
intros b st Hbe contra.
unfold bassn in contra.
rewrite contra in Hbe. inversion Hbe. Qed.

Now we can formalize the Hoare proof rule for conditionals and prove it correct.

Theorem hoare_if : P Q b c1 c2,
{{fun st => P st bassn b st}} c1 {{Q}}
{{fun st => P st ~(bassn b st)}} c2 {{Q}}
{{P}} (IFB b THEN c1 ELSE c2 FI) {{Q}}.
Proof.
intros P Q b c1 c2 HTrue HFalse st st' HE HP.
inversion HE; subst.
Case "b is true".
apply (HTrue st st').
assumption.
split. assumption.
apply bexp_eval_true. assumption.
Case "b is false".
apply (HFalse st st').
assumption.
split. assumption.
apply bexp_eval_false. assumption. Qed.

Here is a formal proof that the program we used to motivate the rule satisfies the specification we gave.

Example if_example :
{{fun st => True}}
IFB (BEq (AId X) (ANum 0))
THEN (Y ::= (ANum 2))
ELSE (Y ::= APlus (AId X) (ANum 1))
FI
{{fun st => asnat (st X) <= asnat (st Y)}}.
Proof.
(* WORKED IN CLASS *)
apply hoare_if.
Case "Then".
eapply hoare_consequence_pre. apply hoare_asgn.
unfold bassn, assn_sub, update. simpl. intros.
inversion H.
symmetry in H1; apply beq_nat_eq in H1.
rewrite H1. omega.
Case "Else".
eapply hoare_consequence_pre. apply hoare_asgn.
unfold assn_sub, update; simpl; intros. omega.
Qed.

### Loops

Finally, we need a rule for reasoning about while loops. Suppose we have a loop
WHILE b DO c END
and we want to find a pre-condition P and a post-condition Q such that
{{P}} WHILE b DO c END {{Q}}
is a valid triple.
First of all, let's think about the case where b is false at the beginning, so that the loop body never executes at all. In this case, the loop behaves like SKIP, so we might be tempted to write
{{P}} WHILE b DO c END {{P}}.
But, as we remarked above for the conditional, we know a little more at the end — not just P, but also the fact that b is false in the current state. So we can enrich the postcondition a little:
{{P}} WHILE b DO c END {{P  ~b}}
What about the case where the loop body does get executed? In order to ensure that P holds when the loop finally exits, we certainly need to make sure that the command c guarantees that P holds whenever c is finished. Moreover, since P holds at the beginning of the first execution of c, and since each execution of c re-establishes P when it finishes, we can always assume that P holds at the beginning of c. This leads us to the following rule:
 {{P}} c {{P}} {{P}} WHILE b DO c END {{P ∧ ~b}}
The proposition P is called an invariant.
This is almost the rule we want, but again it can be improved a little: at the beginning of the loop body, we know not only that P holds, but also that the guard b is true in the current state. This gives us a little more information to use in reasoning about c. Here is the final version of the rule:
 {{P ∧ b}} c {{P}} [hoare_while] {{P}} WHILE b DO c END {{P ∧ ~b}}

Lemma hoare_while : P b c,
{{fun st => P st bassn b st}} c {{P}}
{{P}} WHILE b DO c END {{fun st => P st ~ (bassn b st)}}.
Proof.
intros P b c Hhoare st st' He HP.
(* Like we've seen before, we need to reason by induction
on He, because, in the "keep looping" case, its hypotheses

remember (WHILE b DO c END) as wcom.
ceval_cases (induction He) Case; try (inversion Heqwcom); subst.

Case "E_WhileEnd".
split. assumption. apply bexp_eval_false. assumption.

Case "E_WhileLoop".
apply IHHe2. reflexivity.
apply (Hhoare st st'); try assumption.
split. assumption. apply bexp_eval_true. assumption. Qed.

Example while_example :
{{fun st => asnat (st X) <= 3}}
WHILE (BLe (AId X) (ANum 2))
DO X ::= APlus (AId X) (ANum 1) END
{{fun st => asnat (st X) = 3}}.
Proof.
eapply hoare_consequence_post.
apply hoare_while.
eapply hoare_consequence_pre.
apply hoare_asgn.
unfold bassn, assn_sub. intros. rewrite update_eq. simpl.
inversion H as [_ H0]. simpl in H0. apply ble_nat_true in H0.
omega.
unfold bassn. intros. inversion H as [Hle Hb]. simpl in Hb.
remember (ble_nat (asnat (st X)) 2) as le. destruct le.
apply ex_falso_quodlibet. apply Hb; reflexivity.
symmetry in Heqle. apply ble_nat_false in Heqle. omega.
Qed.

We can also use the while rule to prove the following Hoare triple, which may seem surprising at first...

Theorem always_loop_hoare : P Q,
{{P}} WHILE BTrue DO SKIP END {{Q}}.
Proof.
intros P Q.
apply hoare_consequence_pre with (P' := fun st : state => True).
eapply hoare_consequence_post.
apply hoare_while.
Case "Loop body preserves invariant".
apply hoare_post_true. intros st. apply I.
Case "Loop invariant and negated guard imply postcondition".
simpl. intros st [Hinv Hguard].
apply ex_falso_quodlibet. apply Hguard. reflexivity.
Case "Precondition implies invariant".
intros st H. constructor. Qed.

Actually, this result shouldn't be surprising. If we look back at the definition of hoare_triple, we can see that it asserts something meaningful only when the command terminates.

Print hoare_triple.

If the command doesn't terminate, we can prove anything we like about the post-condition. Here's a more direct proof of the same fact:

Theorem always_loop_hoare' : P Q,
{{P}} WHILE BTrue DO SKIP END {{Q}}.
Proof.
unfold hoare_triple. intros P Q st st' contra.
apply loop_never_stops in contra. inversion contra.
Qed.

Hoare rules that only talk about terminating commands are often said to describe a logic of "partial" correctness. It is also possible to give Hoare rules for "total" correctness, which build in the fact that the commands terminate.

### Exercise: Hoare Rules for REPEAT

Module RepeatExercise.

#### Exercise: 4 stars (hoare_repeat)

In this exercise, we'll add a new constructor to our language of commands: CRepeat. You will write the evaluation rule for repeat and add a new hoare logic theorem to the language for programs involving it.
We recommend that you do this exercise before the ones that follow, as it should help solidify your understanding of the material.

Inductive com : Type :=
| CSkip : com
| CAss : id aexp com
| CSeq : com com com
| CIf : bexp com com com
| CWhile : bexp com com
| CRepeat : com bexp com.

REPEAT behaves like WHILE, except that the loop guard is checked after each execution of the body, with the loop repeating as long as the guard stays false. Because of this, the body will always execute at least once.

Tactic Notation "com_cases" tactic(first) ident(c) :=
first;
[ Case_aux c "SKIP" | Case_aux c "::=" | Case_aux c ";"
| Case_aux c "IFB" | Case_aux c "WHILE" | Case_aux c "CRepeat" ].

Notation "'SKIP'" :=
CSkip.
Notation "c1 ; c2" :=
(CSeq c1 c2) (at level 80, right associativity).
Notation "X '::=' a" :=
(CAss X a) (at level 60).
Notation "'WHILE' b 'DO' c 'END'" :=
(CWhile b c) (at level 80, right associativity).
Notation "'IFB' e1 'THEN' e2 'ELSE' e3 'FI'" :=
(CIf e1 e2 e3) (at level 80, right associativity).
Notation "'REPEAT' e1 'UNTIL' b2 'END'" :=
(CRepeat e1 b2) (at level 80, right associativity).

Add new rules for REPEAT to ceval below. You can use the rules for WHILE as a guide, but remember that the body of a REPEAT should always execute at least once, and that the loop ends when the guard becomes true. Then update the ceval_cases tactic to handle these added cases.

Inductive ceval : state com state Prop :=
| E_Skip : st,
ceval st SKIP st
| E_Ass : st a1 n V,
aeval st a1 = n
ceval st (V ::= a1) (update st V n)
| E_Seq : c1 c2 st st' st'',
ceval st c1 st'
ceval st' c2 st''
ceval st (c1 ; c2) st''
| E_IfTrue : st st' b1 c1 c2,
beval st b1 = true
ceval st c1 st'
ceval st (IFB b1 THEN c1 ELSE c2 FI) st'
| E_IfFalse : st st' b1 c1 c2,
beval st b1 = false
ceval st c2 st'
ceval st (IFB b1 THEN c1 ELSE c2 FI) st'
| E_WhileEnd : b1 st c1,
beval st b1 = false
ceval st (WHILE b1 DO c1 END) st
| E_WhileLoop : st st' st'' b1 c1,
beval st b1 = true
ceval st c1 st'
ceval st' (WHILE b1 DO c1 END) st''
ceval st (WHILE b1 DO c1 END) st''
(* FILL IN HERE *)
.

Tactic Notation "ceval_cases" tactic(first) ident(c) :=
first;
[ Case_aux c "E_Skip" | Case_aux c "E_Ass" | Case_aux c "E_Seq"
| Case_aux c "E_IfTrue" | Case_aux c "E_IfFalse"
| Case_aux c "E_WhileEnd" | Case_aux c "E_WhileLoop"
(* FILL IN HERE *)
].

A couple of definitions from above, copied here so they use the new ceval.

Notation "c1 '/' st '' st'" := (ceval st c1 st')
(at level 40, st at level 39).
Definition hoare_triple (P:Assertion) (c:com) (Q:Assertion)
: Prop :=
st st', (c / st st') P st Q st'.
Notation "{{ P }} c {{ Q }}" := (hoare_triple P c Q) (at level 90, c at next level).

Now state and prove a theorem, hoare_repeat, that expresses an appropriate proof rule for repeat commands. Use hoare_while as a model.

(* FILL IN HERE *)

End RepeatExercise.

## Decorated Programs

The whole point of Hoare Logic is that it is compositional — the structure of proofs exactly follows the structure of programs. This fact suggests that we we could record the essential ideas of a proof informally (leaving out some low-level calculational details) by decorating programs with appropriate assertions around each statement. Such a decorated program carries with it an (informal) proof of its own correctness.
For example, here is a complete decorated program:
{{ True }} =>
{{ x = x }}
X ::= x
{{ X = x }} =>
{{ X = x  z = z }}
Z ::= z;
{{ X = x  Z = z }} =>
{{ Z - X = z - x }}
WHILE X <> 0 DO
{{ Z - X = z - x  X <> 0 }} =>
{{ (Z - 1) - (X - 1) = z - x }}
Z ::= Z - 1;
{{ Z - (X - 1) = z - x }}
X ::= X - 1
{{ Z - X = z - x }}
END;
{{ Z - X = z - x  ~ (X <> 0) }} =>
{{ Z = z - x }} =>
{{ Z + 1 = z - x + 1 }}
Z ::= Z + 1
{{ Z = z - x + 1 }}
Concretely, a decorated program consists of the program text interleaved with assertions. To check that a decorated program represents a valid proof, we check that each individual command is locally consistent with its accompanying assertions in the following sense:
• SKIP is locally consistent if its precondition and postcondition are the same
{{ P }}
SKIP
{{ P }}
• The sequential composition of commands c1 and c2 is locally consistent (with respect to assertions P and R) if c1 is (with respect to P and Q) and c2 is (with respect to Q and R):
{{ P }}
c1;
{{ Q }}
c2
{{ R }}
• An assignment is locally consistent if its precondition is the appropriate substitution of its postcondition:
{{ P where a is substituted for X }}
X ::= a
{{ P }}
• A conditional is locally consistent (with respect to assertions P and Q) if the assertions at the top of its "then" and "else" branches are exactly Pb and P/\~b and if its "then" branch is locally consistent (with respect to Pb and Q) and its "else" branch is locally consistent (with respect to P/\~b and Q):
{{ P }}
IFB b THEN
{{ P  b }}
c1
{{ Q }}
ELSE
{{ P  ~b }}
c2
{{ Q }}
FI
• A while loop is locally consistent if its postcondition is P/\~b (where P is its precondition) and if the pre- and postconditions of its body are exactly Pb and P:
{{ P }}
WHILE b DO
{{ P  b }}
c1
{{ P }}
END
{{ P  ~b }}
• A pair of assertions separated by => is locally consistent if the first implies the second (in all states):
{{ P }} =>
{{ P' }}

# Reasoning About Programs with Hoare Logic

## Example: Slow Subtraction

Informally:
{{ X = x  Z = z }} =>
{{ Z - X = z - x }}
WHILE X <> 0 DO
{{ Z - X = z - x  X <> 0 }} =>
{{ (Z - 1) - (X - 1) = z - x }}
Z ::= Z - 1;
{{ Z - (X - 1) = z - x }}
X ::= X - 1
{{ Z - X = z - x }}
END
{{ Z - X = z - x  ~ (X <> 0) }} =>
{{ Z = z - x }}
Formally:

Definition subtract_slowly : com :=
WHILE BNot (BEq (AId X) (ANum 0)) DO
Z ::= AMinus (AId Z) (ANum 1);
X ::= AMinus (AId X) (ANum 1)
END.

Definition subtract_slowly_invariant x z :=
fun st => minus (asnat (st Z)) (asnat (st X)) = minus z x.

Theorem subtract_slowly_correct : x z,
{{fun st => asnat (st X) = x asnat (st Z) = z}}
subtract_slowly
{{fun st => asnat (st Z) = minus z x}}.
Proof.
(* Note that we do NOT unfold the definition of hoare_triple
anywhere in this proof!  The goal is to use only the hoare
rules.  Your proofs should do the same. *)

intros x z. unfold subtract_slowly.
(* First we need to transform the pre and postconditions so
that hoare_while will apply.  In particular, the
precondition should be the loop invariant. *)

eapply hoare_consequence with (P' := subtract_slowly_invariant x z).
apply hoare_while.

Case "Loop body preserves invariant".
(* Split up the two assignments with hoare_seq - using eapply
lets us solve the second one immediately with hoare_asgn *)

eapply hoare_seq. apply hoare_asgn.
(* Now for the first assignment, transform the precondition
so we can use hoare_asgn *)

eapply hoare_consequence_pre. apply hoare_asgn.
(* Finally, we need to justify the implication generated by
hoare_consequence_pre (this bit of reasoning is elided in the
informal proof) *)

unfold subtract_slowly_invariant, assn_sub, update, bassn. simpl.
intros st [Inv GuardTrue].
(* There are several ways to do the case analysis here...this
one is fairly general: *)

remember (beq_nat (asnat (st X)) 0) as Q; destruct Q.
inversion GuardTrue.
symmetry in HeqQ. apply beq_nat_false in HeqQ.
omega. (* slow but effective! *)
Case "Initial state satisfies invariant".
(* This is the subgoal generated by the precondition part of our
first use of hoare_consequence.  It's the first implication
written in the decorated program (though we elided the actual
proof there). *)

unfold subtract_slowly_invariant.
intros st [HX HZ]. omega.
Case "Invariant and negated guard imply postcondition".
(* This is the subgoal generated by the postcondition part of
out first use of hoare_consequence.  This implication is
the one written after the while loop in the informal proof. *)

intros st [Inv GuardFalse].
unfold subtract_slowly_invariant in Inv.
unfold bassn in GuardFalse. simpl in GuardFalse.
(* Here's a slightly different alternative for the case analysis that
works out well here (but is less general)... *)

destruct (asnat (st X)).
omega.
apply ex_falso_quodlibet. apply GuardFalse. reflexivity.
Qed.

## Exercise: Reduce to Zero

Here is a while loop that is so simple it needs no invariant:
{{ True }}
WHILE X <> 0 DO
{{ True  X <> 0 }} =>
{{ True }}
X ::= X - 1
{{ True }}
END
{{ True  X = 0 }} =>
{{ X = 0 }}
Your job is to translate this proof to Coq. It may help to look at the slow_subtraction proof for ideas.

#### Exercise: 2 stars (reduce_to_zero_correct)

Definition reduce_to_zero : com :=
WHILE BNot (BEq (AId X) (ANum 0)) DO
X ::= AMinus (AId X) (ANum 1)
END.

Theorem reduce_to_zero_correct :
{{fun st => True}}
reduce_to_zero
{{fun st => asnat (st X) = 0}}.
Proof.
(* FILL IN HERE *) Admitted.

The following program adds the variable X into the variable Z by repeatedly decrementing X and incrementing Z.

WHILE BNot (BEq (AId X) (ANum 0)) DO
Z ::= APlus (AId Z) (ANum 1);
X ::= AMinus (AId X) (ANum 1)
END.

Following the pattern of the subtract_slowly example above, pick a precondition and postcondition that give an appropriate specification of add_slowly; then (informally) decorate the program accordingly.

(* FILL IN HERE *)

Now write down your specification of add_slowly formally, as a Coq Hoare_triple, and prove it valid.

(* FILL IN HERE *)

## Example: Parity

Here's another, slightly trickier example. Make sure you understand the decorated program completely. Understanding all the details of the Coq proof is not required, though it is not actually very hard — all the required ideas are already in the informal version.
{{ X = x }} =>
{{ X = x  0 = 0 }}
Y ::= 0;
{{ X = x  Y = 0 }} =>
{{ (Y=0  ev (x-X))  X<=x }}
WHILE X <> 0 DO
{{ (Y=0  ev (x-X))  X<=x  X<>0 }} =>
{{ (1-Y)=0  ev (x-(X-1))  X-1<=x }}
Y ::= 1 - Y;
{{ Y=0  ev (x-(X-1))  X-1<=x }}
X ::= X - 1
{{ Y=0  ev (x-X X<=x }}
END
{{ (Y=0  ev (x-X))  X<=x  ~(X<>0) }} =>
{{ Y=0  ev x }}

Definition find_parity : com :=
Y ::= (ANum 0);
WHILE (BNot (BEq (AId X) (ANum 0))) DO
Y ::= AMinus (ANum 1) (AId Y);
X ::= AMinus (AId X) (ANum 1)
END.

Definition find_parity_invariant x :=
fun st =>
asnat (st X) <= x
(asnat (st Y) = 0 ev (x - asnat (st X)) asnat (st Y) = 1 ~ev (x - asnat (st X))).

(* It turns out that we'll need the following lemma... *)
Lemma not_ev_ev_S_gen: n,
(~ ev n ev (S n))
(~ ev (S n) ev (S (S n))).
Proof.
induction n as [| n'].
Case "n = 0".
split; intros H.
SCase "".
apply ex_falso_quodlibet. apply H. apply ev_0.
SCase "".
apply ev_SS. apply ev_0.
Case "n = S n'".
inversion IHn' as [Hn HSn]. split; intros H.
SCase "".
apply HSn. apply H.
SCase "".
apply ev_SS. apply Hn. intros contra.
apply H. apply ev_SS. apply contra. Qed.

Lemma not_ev_ev_S : n,
~ ev n ev (S n).
Proof.
intros n.
destruct (not_ev_ev_S_gen n) as [H _].
apply H.
Qed.

Theorem find_parity_correct : x,
{{fun st => asnat (st X) = x}}
find_parity
{{fun st => asnat (st Y) = 0 ev x}}.
Proof.
intros x. unfold find_parity.
apply hoare_seq with (Q := find_parity_invariant x).
eapply hoare_consequence.
apply hoare_while with (P := find_parity_invariant x).
Case "Loop body preserves invariant".
eapply hoare_seq.
apply hoare_asgn.
eapply hoare_consequence_pre.
apply hoare_asgn.
intros st [[Inv1 Inv2] GuardTrue].
unfold find_parity_invariant, bassn, assn_sub, aeval in *.
rewrite update_eq.
rewrite (update_neq Y X); auto.
rewrite (update_neq X Y); auto.
rewrite update_eq.
simpl in GuardTrue. destruct (asnat (st X)).
inversion GuardTrue. simpl.
split. omega.
inversion Inv2 as [[H1 H2] | [H1 H2]]; rewrite H1;
[right|left]; (split; simpl; [omega |]).
apply ev_not_ev_S in H2.
replace (S (x - S n)) with (x-n) in H2 by omega.
rewrite minus_n_O. assumption.
apply not_ev_ev_S in H2.
replace (S (x - S n)) with (x - n) in H2 by omega.
rewrite minus_n_O. assumption.
Case "Precondition implies invariant".
intros st H. assumption.
Case "Invariant implies postcondition".
unfold bassn, find_parity_invariant. simpl.
intros st [[Inv1 Inv2] GuardFalse].
destruct (asnat (st X)).
split; intro.
inversion Inv2.
inversion H0 as [_ H1]. replace (x-0) with x in H1 by omega.
assumption.
inversion H0 as [H0' _]. rewrite H in H0'. inversion H0'.
inversion Inv2.
inversion H0. assumption.
inversion H0 as [_ H1]. replace (x-0) with x in H1 by omega.
apply ex_falso_quodlibet. apply H1. assumption.
apply ex_falso_quodlibet. apply GuardFalse. reflexivity.
Case "invariant established before loop".
eapply hoare_consequence_pre.
apply hoare_asgn.
intros st H.
unfold assn_sub, find_parity_invariant, update. simpl.
subst.
split.
omega.
replace (asnat (st X) - asnat (st X)) with 0 by omega.
left. split. reflexivity.
apply ev_0. Qed.

#### Exercise: 3 stars (wrong_find_parity_invariant)

A plausible first attempt at stating an invariant for find_parity is the following.

Definition find_parity_invariant' x :=
fun st =>
(asnat (st Y)) = 0 ev (x - asnat (st X)).

Why doesn't this work? (Hint: Don't waste time trying to answer this exercise by attempting a formal proof and seeing where it goes wrong. Just think about whether the loop body actually preserves the property.)

(* FILL IN HERE *)

## Example: Finding Square Roots

Definition sqrt_loop : com :=
WHILE BLe (AMult (APlus (ANum 1) (AId Z))
(APlus (ANum 1) (AId Z)))
(AId X) DO
Z ::= APlus (ANum 1) (AId Z)
END.

Definition sqrt_com : com :=
Z ::= ANum 0;
sqrt_loop.

Definition sqrt_spec (x : nat) : Assertion :=
fun st =>
((asnat (st Z)) * (asnat (st Z))) <= x
~ (((S (asnat (st Z))) * (S (asnat (st Z)))) <= x).

Definition sqrt_inv (x : nat) : Assertion :=
fun st =>
asnat (st X) = x
((asnat (st Z)) * (asnat (st Z))) <= x.

Theorem random_fact_1 : st,
(S (asnat (st Z))) * (S (asnat (st Z))) <= asnat (st X)
bassn (BLe (AMult (APlus (ANum 1) (AId Z))
(APlus (ANum 1) (AId Z)))
(AId X)) st.
Proof.
intros st Hle. unfold bassn. simpl.
destruct (asnat (st X)) as [|x'].
Case "asnat (st X) = 0".
inversion Hle.
Case "asnat (st X) = S x'".
simpl in Hle. apply le_S_n in Hle.
remember (ble_nat (plus (asnat (st Z))
((asnat (st Z)) * (S (asnat (st Z))))) x')
as ble.
destruct ble. reflexivity.
symmetry in Heqble. apply ble_nat_false in Heqble.
unfold not in Heqble. apply Heqble in Hle. inversion Hle.
Qed.

Theorem random_fact_2 : st,
bassn (BLe (AMult (APlus (ANum 1) (AId Z))
(APlus (ANum 1) (AId Z)))
(AId X)) st
asnat (aeval st (APlus (ANum 1) (AId Z)))
* asnat (aeval st (APlus (ANum 1) (AId Z)))
<= asnat (st X).
Proof.
intros st Hble. unfold bassn in Hble. simpl in *.
destruct (asnat (st X)) as [| x'].
Case "asnat (st X) = 0".
inversion Hble.
Case "asnat (st X) = S x'".
apply ble_nat_true in Hble. omega. Qed.

Theorem sqrt_com_correct : x,
{{fun st => True}} (X ::= ANum x; sqrt_com) {{sqrt_spec x}}.
Proof.
intros x.
apply hoare_seq with (Q := fun st => asnat (st X) = x).
Case "sqrt_com".
unfold sqrt_com.
apply hoare_seq with (Q := fun st => asnat (st X) = x
asnat (st Z) = 0).

SCase "sqrt_loop".
unfold sqrt_loop.
eapply hoare_consequence.
apply hoare_while with (P := sqrt_inv x).

SSCase "loop preserves invariant".
eapply hoare_consequence_pre.
apply hoare_asgn. intros st H.
unfold assn_sub. unfold sqrt_inv in *.
inversion H as [[HX HZ] HP]. split.
SSSCase "X is preserved".
rewrite update_neq; auto.
SSSCase "Z is updated correctly".
rewrite (update_eq (aeval st (APlus (ANum 1) (AId Z))) Z st).
subst. apply random_fact_2. assumption.

SSCase "invariant is true initially".
intros st H. inversion H as [HX HZ].
unfold sqrt_inv. split. assumption.
rewrite HZ. simpl. omega.

SSCase "after loop, spec is satisfied".
intros st H. unfold sqrt_spec.
inversion H as [HX HP].
unfold sqrt_inv in HX. inversion HX as [HX' Harith].
split. assumption.
intros contra. apply HP. clear HP.
simpl. simpl in contra.
apply random_fact_1. subst x. assumption.

SCase "Z set to 0".
eapply hoare_consequence_pre. apply hoare_asgn.
intros st HX.
unfold assn_sub. split.
rewrite update_neq; auto.
rewrite update_eq; auto.

Case "assignment of X".
eapply hoare_consequence_pre. apply hoare_asgn.
intros st H.
unfold assn_sub. rewrite update_eq; auto. Qed.

#### Exercise: 3 stars, optional (sqrt_informal)

Write a decorated program corresponding to the above correctness proof.

(* FILL IN HERE *)

## Exercise: Factorial

Module Factorial.

Fixpoint real_fact (n:nat) : nat :=
match n with
| O => 1
| S n' => n * (real_fact n')
end.

Recall the factorial Imp program:

Definition fact_body : com :=
Y ::= AMult (AId Y) (AId Z);
Z ::= AMinus (AId Z) (ANum 1).

Definition fact_loop : com :=
WHILE BNot (BEq (AId Z) (ANum 0)) DO
fact_body
END.

Definition fact_com : com :=
Z ::= (AId X);
Y ::= ANum 1;
fact_loop.

#### Exercise: 3 stars, optional (fact_informal)

Decorate the fact_com program to show that it satisfies the specification given by the pre and postconditions below. Just as we have done above, you may elide the algebraic reasoning about arithmetic, the less-than relation, etc., that is (formally) required by the rule of consequence:
(* FILL IN HERE *)
{{ X = x }}
Z ::= X;
Y ::= 1;
WHILE Z <> 0 DO
Y ::= Y * Z;
Z ::= Z - 1
END
{{ Y = real_fact x }}

#### Exercise: 4 stars, optional (fact_formal)

Prove formally that fact_com satisfies this specification, using your informal proof as a guide. You may want to state the loop invariant separately (as we did in the examples).

Theorem fact_com_correct : x,
{{fun st => asnat (st X) = x}} fact_com
{{fun st => asnat (st Y) = real_fact x}}.
Proof.
(* FILL IN HERE *) Admitted.

End Factorial.

## Reasoning About Programs with Lists

#### Exercise: 3 stars (list_sum)

Here is a direct definition of the sum of the elements of a list, and an Imp program that computes the sum.

Definition sum l := fold_right plus 0 l.

Definition sum_program :=
Y ::= ANum 0;
WHILE (BIsCons (AId X)) DO
Y ::= APlus (AId Y) (AHead (AId X)) ;
X ::= ATail (AId X)
END.

Provide an informal proof of the following specification of sum_program in the form of a decorated version of the program.

Definition sum_program_spec := l,
{{ fun st => aslist (st X) = l }}
sum_program
{{ fun st => asnat (st Y) = sum l }}.

(* FILL IN HERE *)
Next, let's look at a formal Hoare Logic proof for a program that deals with lists. We will verify the following program, which checks if the number Y occurs in the list X, and if so sets Z to 1.

Definition list_member :=
WHILE BIsCons (AId X) DO
IFB (BEq (AId Y) (AHead (AId X))) THEN
Z ::= (ANum 1)
ELSE
SKIP
FI;
X ::= ATail (AId X)
END.

Definition list_member_spec := l n,
{{ fun st => st X = VList l st Y = VNat n st Z = VNat 0 }}
list_member
{{ fun st => st Z = VNat 1 appears_in n l }}.

The proof we will use, written informally, looks as follows:
{{ X = l  Y = n  Z = 0 }} =>
{{ Y = n   pp ++ X = l  (Z = 1  appears_in n p) }}
WHILE (BIsCons X
DO
{{ Y = n  ( pp ++ X = l  (Z = 1  appears_in n p))
(BIsCons X) }}
{{ Y = n
( pp ++ X = l  (Z = 1  appears_in n p))
(BIsCons X)
Y == AHead X }} =>
{{ Y = n  ( pp ++ tail X = l
(1 = 1  appears_in n p)) }}
Z ::= 1
{{ Y = n  ( pp ++ tail X = l
(Z = 1  appears_in n p)) }}
ELSE
{{ Y = n
( pp ++ X = l  (Z = 1  appears_in n p))
(BIsCons X)
~ (Y == head X) }} =>
{{ Y = n
( pp ++ tail X = l  (Z = 1  appears_in n p)) }}
SKIP
{{ Y = n
( pp ++ tail X = l  (Z = 1  appears_in n p)) }}
FI;
X ::= ATail X
{{ Y = n
( pp ++ X = l  (Z = 1  appears_in n p)) }}
END
{{ Y = n
( pp ++ X = l  (Z = 1  appears_in n p))
~ (BIsCons X) }} =>
{{ Z = 1  appears_in n l }}
The only interesting part of the proof is the choice of loop invariant:
pp ++ X = l  (Z = 1  appears_in n p)
This states that at each iteration of the loop, the original list l is equal to the append of the current value of X and some other list p which is not the value of any variable in the program, but keeps track of enough information from the original state to make the proof go through. (Such a p is sometimes called a "ghost variable").
In order to show that such a list p exists, in each iteration we add the head of X to the end of p. This needs the function snoc, from Poly.v.

Fixpoint snoc {X:Type} (l:list X) (v:X) : (list X) :=
match l with
| nil => [ v ]
| cons h t => h :: (snoc t v)
end.

Lemma snoc_equation : (A : Type) (h:A) (x y : list A),
snoc x h ++ y = x ++ h :: y.
Proof.
intros A h x y.
induction x.
Case "x = []". reflexivity.
Case "x = cons". simpl. rewrite IHx. reflexivity.
Qed.

The main proof uses various lemmas.

Lemma appears_in_snoc1 : a l,
appears_in a (snoc l a).
Proof.
induction l.
Case "l = []". apply ai_here.
Case "l = cons". simpl. apply ai_later. apply IHl.
Qed.

Lemma appears_in_snoc2 : a b l,
appears_in a l
appears_in a (snoc l b).
Proof.
induction l; intros H; inversion H; subst; simpl.
Case "l = []". apply ai_here.
Case "l = cons". apply ai_later. apply IHl. assumption.
Qed.

Lemma appears_in_snoc3 : a b l,
appears_in a (snoc l b)
(appears_in a l a = b).
Proof.
induction l; intros H.
Case "l = []". inversion H.
SCase "ai_here". right. reflexivity.
SCase "ai_later". left. assumption.
Case "l = cons". inversion H; subst.
SCase "ai_here". left. apply ai_here.
SCase "ai_later". destruct (IHl H1).
left. apply ai_later. assumption.
right. assumption.
Qed.

Lemma append_singleton_equation : (x : nat) l l',
(l ++ [x]) ++ l' = l ++ x :: l'.
Proof.
intros x l l'.
induction l.
reflexivity.
simpl. rewrite IHl. reflexivity.
Qed.

Lemma append_nil : (A : Type) (l : list A),
l ++ [] = l.
Proof.
induction l.
reflexivity.
simpl. rewrite IHl. reflexivity.
Qed.

Lemma beq_true__eq : n n',
beq_nat n n' = true
n = n'.
Proof.
induction n; destruct n'.
Case "n = 0, n' = 0". reflexivity.
Case "n = 0, n' = S _". simpl. intros H. inversion H.
Case "n = S, n' = 0". simpl. intros H. inversion H.
Case "n = S, n' = S". simpl. intros H.
rewrite (IHn n' H). reflexivity.
Qed.

Lemma beq_nat_refl : n,
beq_nat n n = true.
Proof.
induction n.
reflexivity.
simpl. assumption.
Qed.

Theorem list_member_correct : l n,
{{ fun st => st X = VList l st Y = VNat n st Z = VNat 0 }}
list_member
{{ fun st => st Z = VNat 1 appears_in n l }}.
Proof.
intros l n.
eapply hoare_consequence.
apply hoare_while with (P := fun st =>
st Y = VNat n
p, p ++ aslist (st X) = l
(st Z = VNat 1 appears_in n p)).
(* The loop body preserves the invariant: *)
eapply hoare_seq.
apply hoare_asgn.
apply hoare_if.
Case "If taken".
eapply hoare_consequence_pre.
apply hoare_asgn.
intros st [[[H1 [p [H2 H3]]] H9] H10].
unfold assn_sub. split.
(* (st Y) is still n *)
rewrite update_neq; try reflexivity.
rewrite update_neq; try reflexivity.
assumption.
(* and the interesting part of the invariant is preserved: *)
(* X has to be a cons *)
remember (aslist (st X)) as x.
destruct x as [|h x'].
unfold bassn in H9. unfold beval in H9. unfold aeval in H9.
rewrite Heqx in H9. inversion H9.

(snoc p h).
rewrite update_eq.
unfold aeval. rewrite update_neq; try reflexivity.
rewrite Heqx.
split.
rewrite snoc_equation. assumption.

rewrite update_neq; try reflexivity.
rewrite update_eq.
split.
simpl.
unfold bassn in H10. unfold beval in H10.
unfold aeval in H10. rewrite H1 in H10.
rewrite Heqx in H10. simpl in H10.
rewrite (beq_true__eq _ _ H10).
intros. apply appears_in_snoc1.

intros. reflexivity.
Case "If not taken".
eapply hoare_consequence_pre. apply hoare_skip.
unfold assn_sub.
intros st [[[H1 [p [H2 H3]]] H9] H10].
split.
(* (st Y) is still n *)
rewrite update_neq; try reflexivity.
assumption.
(* and the interesting part of the invariant is preserved: *)
(* X has to be a cons *)
remember (aslist (st X)) as x.
destruct x as [|h x'].
unfold bassn in H9. unfold beval in H9. unfold aeval in H9.
rewrite Heqx in H9. inversion H9.

(snoc p h).
split.
rewrite update_eq.
unfold aeval. rewrite Heqx.
rewrite snoc_equation. assumption.

rewrite update_neq; try reflexivity.
split.
intros. apply appears_in_snoc2. apply H3. assumption.

intros. destruct (appears_in_snoc3 _ _ _ H).
SCase "later".
inversion H3 as [_ H3'].
apply H3'. assumption.
SCase "here (absurd)".
subst.
unfold bassn in H10. unfold beval in H10. unfold aeval in H10.
rewrite Heqx in H10. rewrite H1 in H10.
simpl in H10. rewrite beq_nat_refl in H10.
apply ex_falso_quodlibet. apply H10. reflexivity.
(* The invariant holds at the start of the loop: *)
intros st [H1 [H2 H3]].
rewrite H1. rewrite H2. rewrite H3.
split.
reflexivity.
[]. split.
reflexivity.
split; intros H; inversion H.
(* At the end of the loop the invariant implies the right thing. *)
simpl. intros st [[H1 [p [H2 H3]]] H5].
(* x must be  *)
unfold bassn in H5. unfold beval in H5. unfold aeval in H5.
destruct (aslist (st X)) as [|h x'].
rewrite append_nil in H2.
rewrite H2.
assumption.

apply ex_falso_quodlibet. apply H5. reflexivity.
Qed.

#### Exercise: 4 stars, optional (list_reverse)

Recall the function rev from Poly.v, for reversing lists.

Fixpoint rev {X:Type} (l:list X) : list X :=
match l with
| nil => []
| cons h t => snoc (rev t) h
end.

Write an Imp program list_reverse_program that reverses lists. Formally prove that it satisfies the following specification:
l : list nat,
{{ X =  l  Y = nil }}
list_reverse_program
{{ Y = rev l }}.
You may find the lemmas append_nil and rev_equation useful.

Lemma rev_equation : (A : Type) (h : A) (x y : list A),
rev (h :: x) ++ y = rev x ++ h :: y.
Proof.
intros. simpl. apply snoc_equation.
Qed.

(* FILL IN HERE *)

# Formalizing Decorated Programs

The informal conventions for decorated programs amount to a way of displaying Hoare triples in which commands are annotated with enough embedded assertions that checking the validity of the triple is reduced to simple algebraic calculations showing that some assertions were stronger than others.
In this section, we show that this informal presentation style can actually be made completely formal.

## Syntax

The first thing we need to do is to formalize a variant of the syntax of commands that includes embedded assertions. We call the new commands decorated commands, or dcoms.

Inductive dcom : Type :=
| DCSkip : Assertion dcom
| DCSeq : dcom dcom dcom
| DCAsgn : id aexp Assertion dcom
| DCIf : bexp Assertion dcom Assertion dcom dcom
| DCWhile : bexp Assertion dcom Assertion dcom
| DCPre : Assertion dcom dcom
| DCPost : dcom Assertion dcom.

Tactic Notation "dcom_cases" tactic(first) ident(c) :=
first;
[ Case_aux c "Skip" | Case_aux c "Seq" | Case_aux c "Asgn"
| Case_aux c "If" | Case_aux c "While"
| Case_aux c "Pre" | Case_aux c "Post" ].

Notation "'SKIP' {{ P }}"
:= (DCSkip P)
(at level 10) : dcom_scope.
Notation "l '::=' a {{ P }}"
:= (DCAsgn l a P)
(at level 60, a at next level) : dcom_scope.
Notation "'WHILE' b 'DO' {{ Pbody }} d 'END' {{ Ppost }}"
:= (DCWhile b Pbody d Ppost)
(at level 80, right associativity) : dcom_scope.
Notation "'IFB' b 'THEN' {{ P }} d 'ELSE' {{ P' }} d' 'FI'"
:= (DCIf b P d P' d')
(at level 80, right associativity) : dcom_scope.
Notation "'=>' {{ P }} d"
:= (DCPre P d)
(at level 90, right associativity) : dcom_scope.
Notation "{{ P }} d"
:= (DCPre P d)
(at level 90) : dcom_scope.
Notation "d '=>' {{ P }}"
:= (DCPost d P)
(at level 91, right associativity) : dcom_scope.
Notation " d ; d' "
:= (DCSeq d d')
(at level 80, right associativity) : dcom_scope.

Delimit Scope dcom_scope with dcom.

To avoid clashing with the existing Notation definitions for ordinary commands, we introduce these notations in a special scope called dcom_scope, and we wrap examples with the declaration % dcom to signal that we want the notations to be interpreted in this scope.
Careful readers will note that we've defined two notations for the DCPre constructor, one with and one without a =>. The "without" version is intended to be used to supply the initial precondition at the very top of the program.

Example dec_while : dcom := (
{{ fun st => True }}
WHILE (BNot (BEq (AId X) (ANum 0)))
DO
{{ fun st => ~(asnat (st X) = 0) }}
X ::= (AMinus (AId X) (ANum 1))
{{ fun _ => True }}
END
{{ fun st => asnat (st X) = 0 }}
) % dcom.

It is easy to go from a dcom to a com by erasing all annotations.

Fixpoint extract (d:dcom) : com :=
match d with
| DCSkip _ => SKIP
| DCSeq d1 d2 => (extract d1 ; extract d2)
| DCAsgn V a _ => V ::= a
| DCIf b _ d1 _ d2 => IFB b THEN extract d1 ELSE extract d2 FI
| DCWhile b _ d _ => WHILE b DO extract d END
| DCPre _ d => extract d
| DCPost d _ => extract d
end.

The choice of exactly where to put assertions in the definition of dcom is a bit subtle. The simplest thing to do would be to annotate every dcom with a precondition and postcondition. But this would result in very verbose programs with a lot of repeated annotations: for example, a program like SKIP;SKIP would have to be annotated as
{{P}} ({{P}} SKIP {{P}}) ; ({{P}} SKIP {{P}}) {{P}},
with pre- and post-conditions on each SKIP, plus identical pre- and post-conditions on the semicolon!
Instead, the rule we've followed is this:
• The post-condition expected by each dcom d is embedded in d
• The pre-condition is supplied by the context.
In other words, the invariant of the representation is that a dcom d together with a precondition P determines a Hoare triple {{P}} (extract d) {{post d}}, where post is defined as follows:

Fixpoint post (d:dcom) : Assertion :=
match d with
| DCSkip P => P
| DCSeq d1 d2 => post d2
| DCAsgn V a Q => Q
| DCIf _ _ d1 _ d2 => post d1
| DCWhile b Pbody c Ppost => Ppost
| DCPre _ d => post d
| DCPost c Q => Q
end.

We can define a similar function that extracts the "initial precondition" from a decorated program.

Fixpoint pre (d:dcom) : Assertion :=
match d with
| DCSkip P => fun st => True
| DCSeq c1 c2 => pre c1
| DCAsgn V a Q => fun st => True
| DCIf _ _ t _ e => fun st => True
| DCWhile b Pbody c Ppost => fun st => True
| DCPre P c => P
| DCPost c Q => pre c
end.

This function is not doing anything sophisticated like calculating a weakest precondition; it just recursively searches for an explicit annotation at the very beginning of the program, returning default answers for programs that lack an explicit precondition (like a bare assignment or SKIP).
Using pre and post, and assuming that we adopt the convention of always supplying an explicit precondition annotation at the very beginning of our decorated programs, we can express what it means for a decorated program to be correct as follows:

Definition dec_correct (d:dcom) :=
{{pre d}} (extract d) {{post d}}.

To check whether this Hoare triple is valid, we need a way to extract the "proof obligations" from a decorated program. These obligations are often called verification conditions, because they are the facts that must be verified (by some process looking at the decorated program) to see that the decorations are logically consistent and thus add up to a proof of correctness.

## Extracting Verification Conditions

First, a bit of notation:

Definition assert_implies (P Q : Assertion) : Prop :=
st, P st Q st.

We will write P Q (in ASCII, P ~~> Q) for assert_implies P Q.

Notation "P Q" := (assert_implies P Q) (at level 80).
Notation "P Q" := (P Q Q P) (at level 80).

Next, the key definition. The function verification_conditions takes a dcom d together with a precondition P and returns a proposition that, if it can be proved, implies that the triple {{P}} (extract d) {{post d}} is valid. It does this by walking over d and generating a big conjunction including all the "local checks" that we listed when we described the informal rules for decorated programs. (Strictly speaking, we need to massage the informal rules a little bit to add some uses of the rule of consequence, but the correspondence should be clear.)

Fixpoint verification_conditions (P : Assertion) (d:dcom) : Prop :=
match d with
| DCSkip Q =>
(P Q)
| DCSeq d1 d2 =>
verification_conditions P d1
verification_conditions (post d1) d2
| DCAsgn V a Q =>
(P assn_sub V a Q)
| DCIf b P1 t P2 e =>
((fun st => P st bassn b st) P1)
((fun st => P st ~ (bassn b st)) P2)
(post t = post e)
verification_conditions P1 t
verification_conditions P2 e
| DCWhile b Pbody d Ppost =>
(* post d is the loop invariant and the initial precondition *)
(P post d)
((fun st => post d st bassn b st) Pbody)
((fun st => post d st ~(bassn b st)) Ppost)
verification_conditions (fun st => post d st bassn b st) d
| DCPre P' d =>
(P P') verification_conditions P' d
| DCPost d Q =>
verification_conditions P d (post d Q)
end.

And now, the key theorem, which captures the claim that the verification_conditions function does its job correctly. Not surprisingly, we need all of the Hoare Logic rules in the proof. We have used in variants of several tactics before to apply them to values in the context rather than the goal. An extension of this idea is the syntax tactic in *, which applies tactic in the goal and every hypothesis in the context. We most commonly use this facility in conjunction with the simpl tactic, as below.

Theorem verification_correct : d P,
verification_conditions P d {{P}} (extract d) {{post d}}.
Proof.
dcom_cases (induction d) Case; intros P H; simpl in *.
Case "Skip".
eapply hoare_consequence_pre.
apply hoare_skip.
assumption.
Case "Seq".
inversion H as [H1 H2]. clear H.
eapply hoare_seq.
apply IHd2. apply H2.
apply IHd1. apply H1.
Case "Asgn".
eapply hoare_consequence_pre.
apply hoare_asgn.
assumption.
Case "If".
inversion H as [HPre1 [HPre2 [HQeq [HThen HElse]]]]; clear H.
apply hoare_if.
eapply hoare_consequence_pre. apply IHd1. apply HThen. assumption.
simpl. rewrite HQeq.
eapply hoare_consequence_pre. apply IHd2. apply HElse. assumption.
Case "While".
rename a into Pbody. rename a0 into Ppost.
inversion H as [Hpre [Hbody [Hpost Hd]]]; clear H.
eapply hoare_consequence.
apply hoare_while with (P := post d).
apply IHd. apply Hd.
assumption. apply Hpost.
Case "Pre".
inversion H as [HP Hd]; clear H.
eapply hoare_consequence_pre. apply IHd. apply Hd. assumption.
Case "Post".
inversion H as [Hd HQ]; clear H.
eapply hoare_consequence_post. apply IHd. apply Hd. assumption.
Qed.

## Examples

The propositions generated by verification_conditions are fairly big, and they contain many conjuncts that are essentially trivial.

Eval simpl in (verification_conditions (fun st => True) dec_while).
(* ====>
((fun _ : state => True) ~~> (fun _ : state => True)) /\
((fun _ : state => True) ~~> (fun _ : state => True)) /\
((fun st : state => True /\ bassn (BNot (BEq (AId X) (ANum 0))) st)
<~~> (fun st : state => asnat (st X) <> 0)) /\
((fun st : state => True /\ ~ bassn (BNot (BEq (AId X) (ANum 0))) st)
<~~> (fun st : state => asnat (st X) = 0)) /\
(fun st : state => True /\ bassn (BNot (BEq (AId X) (ANum 0))) st)
~~> assn_sub X (AMinus (AId X) (ANum 1)) (fun _ : state => True) *)

We can certainly work with them using just the tactics we have so far, but we can make things much smoother with a bit of automation. We first define a custom verify tactic that applies splitting repeatedly to turn all the conjunctions into separate subgoals and then uses omega and eauto (a handy general-purpose automation tactic that we'll discuss in detail later) to deal with as many of them as possible.

Tactic Notation "verify" :=
try apply verification_correct;
repeat split;
simpl; unfold assert_implies;
unfold bassn in *; unfold beval in *; unfold aeval in *;
unfold assn_sub; simpl in *;
intros;
repeat match goal with [H : _ _ _] => destruct H end;
try eauto; try omega.

What's left after verify does its thing is "just the interesting parts" of checking that the decorations are correct. For example:

Theorem dec_while_correct :
dec_correct dec_while.
Proof.
verify; destruct (asnat (st X)).
inversion H0.
unfold not. intros. inversion H1.
apply ex_falso_quodlibet. apply H. reflexivity.
reflexivity.
reflexivity.
apply ex_falso_quodlibet. apply H0. reflexivity.
unfold not. intros. inversion H0.
inversion H.
Qed.

Another example (formalizing a decorated program we've seen before):

Example subtract_slowly_dec (x:nat) (z:nat) : dcom := (
{{ fun st => asnat (st X) = x asnat (st Z) = z }}
WHILE BNot (BEq (AId X) (ANum 0))
DO {{ fun st => asnat (st Z) - asnat (st X) = z - x
bassn (BNot (BEq (AId X) (ANum 0))) st }}
Z ::= AMinus (AId Z) (ANum 1)
{{ fun st => asnat (st Z) - (asnat (st X) - 1) = z - x }} ;
X ::= AMinus (AId X) (ANum 1)
{{ fun st => asnat (st Z) - asnat (st X) = z - x }}
END
{{ fun st => asnat (st Z)
- asnat (st X)
= z - x
~ bassn (BNot (BEq (AId X) (ANum 0))) st }}
=>
{{ fun st => asnat (st Z) = z - x }}
) % dcom.

Theorem subtract_slowly_dec_correct : x z,
dec_correct (subtract_slowly_dec x z).
Proof.
intros. verify.
rewrite H.
assert (H1: update st Z (VNat (asnat (st Z) - 1)) X = st X).
apply update_neq. reflexivity.
rewrite H1. destruct (asnat (st X)) as [| X'].
inversion H0. simpl. rewrite minus_n_O. omega.
destruct (asnat (st X)).
omega.
apply ex_falso_quodlibet. apply H0. reflexivity.
Qed.

#### Exercise: 3 stars (slow_assignment_dec)

A roundabout way of assigning a number currently stored in X to the variable Y is to start Y at 0, then decrement X until it hits 0, incrementing Y at each step.
Here is an informal decorated program that implements this idea given a parameter x:
{{ True }}
X ::= x
{{ X = x }} ;
Y ::= 0
{{ X = x  Y = 0 }} ;
WHILE X <> 0 DO
{{ X + Y = x  X > 0 }}
X ::= X - 1
{{ Y + X + 1 = x }} ;
Y ::= Y + 1
{{ Y + X = x }}
END
{{ Y = x  X = 0 }}
Write a corresponding function that returns a value of type dcom and prove it correct.

(* FILL IN HERE *)

#### Exercise: 4 stars, optional (factorial_dec)

Remember the factorial function we worked with before:

Fixpoint real_fact (n:nat) : nat :=
match n with
| O => 1
| S n' => n * (real_fact n')
end.

Following the pattern of subtract_slowly_dec, write a decorated Imp program that implements the factorial function, and prove it correct.

(* FILL IN HERE *)
Finally, for a bigger example, let's redo the proof of list_member_correct from above using our new tools.
Notice that the verify tactic leaves subgoals for each use of hoare_consequence — that is, for each => that occurs in the decorated program. Each of these implications relies on a fact about lists, for example that l ++ [] = l. In other words, the Hoare logic infrastructure has taken care of the boilerplate reasoning about the execution of imperative programs, while the user has to prove lemmas that are specific to the problem domain (e.g. lists or numbers).

Definition list_member_dec (n : nat) (l : list nat) : dcom := (
{{ fun st => st X = VList l st Y = VNat n st Z = VNat 0 }}
WHILE BIsCons (AId X)
DO {{ fun st => st Y = VNat n
( p, p ++ aslist (st X) = l
(st Z = VNat 1 appears_in n p))
bassn (BIsCons (AId X)) st }}
IFB (BEq (AId Y) (AHead (AId X))) THEN
{{ fun st =>
((st Y = VNat n
( p, p ++ aslist (st X) = l
(st Z = VNat 1 appears_in n p)))
bassn (BIsCons (AId X)) st)
bassn (BEq (AId Y) (AHead (AId X))) st }}
=>
{{ fun st =>
st Y = VNat n
( p, p ++ tail (aslist (st X)) = l
(VNat 1 = VNat 1 appears_in n p)) }}
Z ::= ANum 1
{{ fun st => st Y = VNat n
( p, p ++ tail (aslist (st X)) = l
(st Z = VNat 1 appears_in n p)) }}
ELSE
{{ fun st =>
((st Y = VNat n
( p, p ++ aslist (st X) = l
(st Z = VNat 1 appears_in n p)))
bassn (BIsCons (AId X)) st)
~bassn (BEq (AId Y) (AHead (AId X))) st }}
=>
{{ fun st =>
st Y = VNat n
( p, p ++ tail (aslist (st X)) = l
(st Z = VNat 1 appears_in n p)) }}
SKIP
{{ fun st => st Y = VNat n
( p, p ++ tail (aslist (st X)) = l
(st Z = VNat 1 appears_in n p)) }}
FI ;
X ::= (ATail (AId X))
{{ fun st =>
st Y = VNat n
( p : list nat, p ++ aslist (st X) = l
(st Z = VNat 1 appears_in n p)) }}
END
{{ fun st =>
(st Y = VNat n
( p, p ++ aslist (st X) = l
(st Z = VNat 1 appears_in n p)))
~bassn (BIsCons (AId X)) st }}
=>
{{ fun st => st Z = VNat 1 appears_in n l }}
) %dcom.

Theorem list_member_correct' : n l,
dec_correct (list_member_dec n l).
Proof.
intros n l.
verify.
Case "The loop precondition holds.".
[]. simpl. split.
rewrite H. reflexivity.
rewrite H1. split; inversion 1.
Case "IF taken".
destruct H2 as [p [H3 H4]].
(* We know X is non-nil. *)
remember (aslist (st X)) as x.
destruct x as [|h x'].
inversion H1.
(snoc p h).
simpl. split.
rewrite snoc_equation. assumption.
split.
rewrite H in H0.
simpl in H0.
rewrite (beq_true__eq _ _ H0).
intros. apply appears_in_snoc1.
intros. reflexivity.
Case "If not taken".
destruct H2 as [p [H3 H4]].
(* We know X is non-nil. *)
remember (aslist (st X)) as x.
destruct x as [|h x'].
inversion H1.
(snoc p h).
split.
rewrite snoc_equation. assumption.
split.
intros. apply appears_in_snoc2. apply H4. assumption.
intros Hai. destruct (appears_in_snoc3 _ _ _ Hai).
SCase "later". apply H4. assumption.
SCase "here (absurd)".
subst.
simpl in H0. rewrite H in H0. rewrite beq_nat_refl in H0.
apply ex_falso_quodlibet. apply H0. reflexivity.
Case "Loop postcondition implies desired conclusion (->)".
destruct H2 as [p [H3 H4]].
unfold bassn in H1. simpl in H1.
destruct (aslist (st X)) as [|h x'].
rewrite append_nil in H3. subst. apply H4. assumption.
apply ex_falso_quodlibet. apply H1. reflexivity.
Case "loop postcondition implies desired conclusion (<-)".
destruct H2 as [p [H3 H4]].
unfold bassn in H1. simpl in H1.
destruct (aslist (st X)) as [|h x'].
rewrite append_nil in H3. subst. apply H4. assumption.
apply ex_falso_quodlibet. apply H1. reflexivity.
Qed.