# RulesReasoning Rules for Term Constructs

Set Implicit Arguments.

From SLF Require Export LibSepReference.
From SLF Require Basic.

Implicit Types n : int.
Implicit Types t : trm.
Implicit Types r v : val.
Implicit Types p : loc.
Implicit Types H : hprop.
Implicit Types Q : valhprop.

# First Pass

The present chapter establishes the remaining reasoning rules of Separation Logic: rules for reasoning about each language construct, and rules for reasoning about primitive operations. All these reasoning rules are expressed using the predicate triple, and are proved with respect to the formal semantics captured by the predicate eval introduced in the previous chapter (Triples). Throughout the chapter, one needs to have in mind the definition of triple:

Definition triple (t:trm) (H:hprop) (Q:valhprop) : Prop :=
s, H seval s t Q.
After establishing the reasoning rules, the present chapter presents an example proof of a concrete program, using directly these reasoning rules, as opposed to using high-level "x-tactics" like done in the first two chapters. This example proof reveals the proof patterns employed by the x-tactics. It also motivates the introduction of the x-tactics to avoid the tedious, repetitive proof patterns.

## Rules for Term Constructs

### Reasoning Rule for Sequences

Let us begin with the reasoning rule for sequences. The Separation Logic reasoning rule for a sequence t1;t2 is essentially the same as that from Hoare logic. The rule is:
 {H} t1 {fun v => H1}     {H1} t2 {Q} {H} (t1;t2) {Q}
The Coq statement corresponding to the above rule is:
Lemma triple_seq : t1 t2 H Q H1,
triple t1 H (fun vH1)
triple t2 H1 Q
triple (trm_seq t1 t2) H Q.
The variable v denotes the result of the evaluation of t1. It is unconstrained in the rule. For well-typed programs, this result is necessarily the unit value. For untyped programs, the result of t1 is simply irrelevant.

#### Exercise: 2 stars, standard, especially useful (triple_seq)

Prove triple_seq by unfolding triple and using eval_seq.
Proof using. (* FILL IN HERE *) Admitted.

### Reasoning Rule for Let-Bindings

Next, we present the reasoning rule for let-bindings. Here again, there is nothing specific to Separation Logic. The rule is exactly the same as in Hoare Logic. The reasoning rule for a let binding let x = t1 in t2 could be stated, in informal writing, in the form:
 {H} t1 {Q1}     (forall x, {Q1 x} t2 {Q}) {H} (let x = t1 in t2) {Q}
Yet, such a presentation makes a confusion between the x that denotes a program variable in let x = t1 in t2, and the x that denotes a universally quantified Coq. The correct statement involves a substitution from the variable x to a value quantified as (v:val).
 {H} t1 {Q1}     (forall v, {Q1 v} (subst x v t2) {Q}) {H} (let x = t1 in t2) {Q}
The corresponding Coq statement is thus as follows.
Lemma triple_let : x t1 t2 Q1 H Q,
triple t1 H Q1
( v1, triple (subst x v1 t2) (Q1 v1) Q)
triple (trm_let x t1 t2) H Q.
Proof using. introv M1 M2 Hs. applys* eval_let. Qed.

### Reasoning Rule for Conditionals

The rule for a conditional is, again, exactly like in Hoare logic.
 b = true -> {H} t1 {Q}     b = false -> {H} t2 {Q} {H} (if b then t1 in t2) {Q}
The corresponding Coq statement appears next.

#### Exercise: 2 stars, standard, especially useful (triple_if_case)

Prove triple_if_case by unfolding triple and using eval_if. Hint: use the tactic case_if to perform a case analysis.
Proof using. (* FILL IN HERE *) Admitted.
The two premises of the rule for if-statements may be factorized into a single one using Coq's conditional construct, like in eval_if.
Lemma triple_if : (b:bool) t1 t2 H Q,
triple (if b then t1 else t2) H Q
triple (trm_if b t1 t2) H Q.
Proof using. introv M Hs. applys* eval_if. Qed.

### Reasoning Rule for Values

The triple specifying the behavior of a value v can be written as a triple with an empty precondition and a postcondition asserting that the result value r is equal to v, in the empty heap. Formally:
 {\[]} v {fun r => \[r = v]}
In practice, however, it is more convenient in practice to work with a judgment whose conclusion is of the form {H} v {Q}, for an arbitrary precondition H and an arbitrary postcondition Q.
 H ==> Q v {H} v {Q}
The Coq statement of the rule for values is thus as follows.
Lemma triple_val : v H Q,
H ==> Q v
triple (trm_val v) H Q.
Proof using. introv M Hs. applys* eval_val. Qed.
It may not be completely obvious at first sight why triple_val is equivalent to the "minimalistic" rule {\[]} v {fun r \[r = v]}. Let us prove the equivalence.

#### Exercise: 1 star, standard, especially useful (triple_val_minimal)

Prove that the alternative rule for values derivable from triple_val. Hint: use the tactic xsimpl to conclude the proof.
Lemma triple_val_minimal : v,
triple (trm_val v) \[] (fun r\[r = v]).
Proof using. (* FILL IN HERE *) Admitted.

#### Exercise: 2 stars, standard, especially useful (triple_val')

More interestingly, prove that triple_val is derivable from triple_val_minimal. Concretely, your goal is to prove triple_val' without unfolding the definition of triple. Use xsimpl, triple_val_minimal and the structural rules on triple such as triple_conseq_frame.
Lemma triple_val' : v H Q,
H ==> Q v
triple (trm_val v) H Q.
Proof using. (* FILL IN HERE *) Admitted.

#### Exercise: 4 stars, standard, especially useful (triple_let_val)

Consider a term of the form let x = v1 in t2, that is, where the argument of the let-binding is already a value. State and prove a reasoning rule of the form:
Lemma triple_let_val : x v1 t2 H Q,
... →
triple (trm_let x v1 t2) H Q.
(* FILL IN HERE *)

### Reasoning Rule for Function Definitions

In addition to the reasoning rule for values, we need reasoning rules for functions and recursive functions that appear as terms in the source program (as opposed to appearing as values).
A function definition trm_fun x t1, expressed as a subterm in a program, evaluates to a value, more precisely to val_fun x t1. Again, we could consider a rule with an empty precondition:
 {\[]} (trm_fun x t1) {fun r => \[r = val_fun x t1]}
However, it is easier to apply a rule featuring a conclusion of the form {H} (trm_fun x t1) {Q}. We thus consider the following rule, very similar to triple_val.
Lemma triple_fun : x t1 H Q,
H ==> Q (val_fun x t1)
triple (trm_fun x t1) H Q.
Proof using. introv M Hs. applys* eval_fun. Qed.
This reasoning rules for functions generalizes to recursive functions. A term describing a recursive function is written trm_fix f x t1, and the corresponding value is written val_fix f x t1.
Lemma triple_fix : f x t1 H Q,
H ==> Q (val_fix f x t1)
triple (trm_fix f x t1) H Q.
Proof using. introv M Hs. applys* eval_fix. Qed.

### Reasoning Rule for Function Calls

Last but not least, we need a reasoning rule to reason about a function application. Consider an application trm_app v1 v2. Assume v1 to be a function, that is, to be of the form val_fun x t1. Then, according to the beta-reduction rule, the semantics of trm_app v1 v2 is the same as that of subst x v2 t1. This reasoning rule is thus:
 v1 = val_fun x t1     {H} (subst x v2 t1) {Q} {H} (trm_app v1 v2) {Q}
The corresponding Coq statement is as shown below.
Lemma triple_app_fun : x v1 v2 t1 H Q,
v1 = val_fun x t1
triple (subst x v2 t1) H Q
triple (trm_app v1 v2) H Q.
Proof using. introv E M Hs. applys* eval_app_fun. Qed.
The reasoning rule that corresponds to beta-reduction for a recursive function involves two substitutions: a first substitution for recursive occurrences of the function, followed with a second substitution for the argument provided to the call.
Lemma triple_app_fix : v1 v2 f x t1 H Q,
v1 = val_fix f x t1
triple (subst x v2 (subst f v1 t1)) H Q
triple (trm_app v1 v2) H Q.
Proof using. introv E M Hs. applys* eval_app_fix. Qed.
This concludes the presentation of the reasoning rules for term constructs. Before we can tackle the verification of actual programs, there remains to present the specifications for the primitive operations. We start with arithmetic operations, then consider heap-manipulating operations.

## Rules for Primitive Arithmetic Operations

Consider a term of the form val_add n1 n2, which is short for trm_app (trm_app (trm_val val_add) (val_int n1)) (val_int n2), if we unfold the coercion val_int.
The addition operation may execute in an empty state. It does not modify the state, and returns the value val_int (n1+n2). Recall the statement of the evaluation rule, in omni-big-step style.
Parameter eval_add : s n1 n2 Q,
Q (val_int (n1 + n2)) s
eval s (val_add (val_int n1) (val_int n2)) Q.
In the specification shown below, the precondition is written \[] and the postcondition binds a return value r of type val specified to be equal to val_int (n1+n2).
Lemma triple_add : n1 n2,
triple (val_add n1 n2)
\[]
(fun r\[r = val_int (n1 + n2)]).
Proof using.
introv Hs. applys* eval_add. lets ->: hempty_inv Hs.
applys hpure_intro. auto.
Qed.

### Division

The specification of the division operation val_div n1 n2 is similar, yet with the extra requirement that the argument n2 must be nonzero. This requirement n2 0 appears in the evaluation rule as well as in the reasoning rule.
Parameter eval_div : s n1 n2 Q,
n2 0
Q (val_int (Z.quot n1 n2)) s
eval s (val_div (val_int n1) (val_int n2)) Q.

Lemma triple_div : n1 n2,
n2 0
triple (val_div n1 n2)
\[]
(fun r\[r = val_int (Z.quot n1 n2)]).
Proof using.
introv Hn2 Hs. applys* eval_div. lets ->: hempty_inv Hs.
applys hpure_intro. auto.
Qed.

### Random-Number Generation

The evaluation of val_rand n is particularly interesting because it involves nondeterminism. Assume n > 0. The output value may be any n1 in the range 0 n1 < n. The statement eval s (val_rand n) Q should capture the fact that Q s n1 holds, for any possible output integer n1. This condition is captured by n1, 0 n1 < n Q n1 s.
Parameter eval_rand : s n Q,
n > 0
( n1, 0 n1 < n Q n1 s)
eval s (val_rand (val_int n)) Q.
The postcondition of the triple for val_rand n asserts that the output value r corresponds to an integer n1, in the range from 0 n1 < n.

#### Exercise: 2 stars, standard, especially useful (triple_rand)

Prove the reasoning rule for calls to the random number generator.
Lemma triple_rand : n,
n > 0
triple (val_rand n)
\[]
(fun r\[ n1, r = val_int n1 0 n1 < n]).
Proof using. (* FILL IN HERE *) Admitted.

## Rules for Primitive Heap-Manipulating Operations

Establishing triples for the operations ref, get, set and free require reasoning about the definition of certain primitive heap predicates. This reasoning is to be performed using the introduction and elimination lemmas introduced in chapter Hprop. For example, hsingle_intro proves (p ~~> v) (Fmap.single p v). Besides, the lemma hstar_hpure_l establishing (\[P] \* H) h = (P H h) will be used pervasively.

### Allocation

Recall that val_ref v allocates a cell with contents v. A call to val_ref v does not depend on the current contents of the state. The operation extends the state with a fresh singleton cell, at some location p, assigning it v as contents. This fresh cell is described by the heap predicate p ~~> v. The evaluation of val_ref v produces the value val_loc p. Thus, if r denotes the result value, we have r = val_loc p for some p. In the corresponding specification shown below, observe how the location p is existentially quantified in the postcondition.
Lemma triple_ref : v,
triple (val_ref v)
\[]
(fun r\ p, \[r = val_loc p] \* p ~~> v).
Proof using.
intros. intros s1 K. applys eval_ref. intros p D.
lets ->: hempty_inv K. rewrite Fmap.update_empty.
applys hexists_intro p. rewrite hstar_hpure_l. split*.
applys hsingle_intro.
Qed.
Using the notation funloc p H as a shorthand for fun (r:val) \ (p:loc), \[r = val_loc p] \* H, the specification for val_ref becomes more concise.
Lemma triple_ref' : v,
triple (val_ref v)
\[]
(funloc p p ~~> v).
Proof using. apply triple_ref. Qed.

### Deallocation

Recall that val_free denotes the operation for deallocating a cell at a given address. A call of the form val_free p executes safely in a state described by p ~~> v. The operation leaves an empty state, and asserts that the return value, named r, is equal to unit.

#### Exercise: 2 stars, standard, especially useful (triple_free')

Prove the reasoning rule for deallocation of a single cell. Hint: use Fmap.indom_single and Fmap.remove_single.
Lemma triple_free' : p v,
triple (val_free (val_loc p))
(p ~~> v)
(fun r\[r = val_unit]).
Proof using. (* FILL IN HERE *) Admitted.
In practice, the information r = val_unit is generally useless. Thus, we state triple_free with an empty postcondition.
Lemma triple_free : p v,
triple (val_free (val_loc p))
(p ~~> v)
(fun _\[]).
Proof using. intros. applys triple_conseq triple_free'; xsimpl*. Qed.

Recall that val_get denotes the operation for reading a memory cell. A call of the form val_get v' executes safely if its argument v' is a value of the form val_loc p for some location p, in a state that features a memory cell at location p, storing some contents v. Such a state is described as p ~~> v. The read operation returns a value r such that r = v, and the memory state of the operation remains unchanged. The specification of val_get is thus expressed as follows.
Lemma triple_get : v p,
triple (val_get p)
(p ~~> v)
(fun r\[r = v] \* (p ~~> v)).
Proof using.
intros. intros s K. lets ->: hsingle_inv K. applys eval_get.
{ applys* Fmap.indom_single. }
{ rewrite hstar_hpure_l. split*. rewrite* Fmap.read_single. }
Qed.

### Write

Recall that val_set denotes the operation for writing a memory cell. A call of the form val_set v' w executes safely if v' is of the form val_loc p for some location p, in a state p ~~> v. The write operation updates this state to p ~~> w, and returns the unit value, which can be ignored. Hence, val_set is specified as follows.

#### Exercise: 2 stars, standard, especially useful (triple_set)

Prove the reasoning rule for a write operation. Hint: use Fmap.update_single.
Lemma triple_set : w p v,
triple (val_set (val_loc p) v)
(p ~~> w)
(fun r\[r = val_unit] \* (p ~~> v)).
Proof using. (* FILL IN HERE *) Admitted.

## Program Verification using the Reasoning Rules of Separation Logic

We have at hand all the necessary rules for carrying out actual verification proofs in Separation Logic.
Module ExamplePrograms.
Export ProgramSyntax.

### Proof of incr

First, we consider the verification of the increment function, which is written in OCaml syntax as:
let incr p =
p := !p + 1
Recall that, for simplicity, we assume programs to be written in "A-normal form", that is, with all intermediate expressions named by a let-binding. Thereafter, we thus consider the following definition for the incr.
let incr p =
let n = !p in
let m = n+1 in
p := m
Using the construct from our embedded programming language, the definition of incr is formalized as shown below.
Definition incr : val :=
val_fun "p" (
trm_let "n" (trm_app val_get (trm_var "p")) (
trm_let "m" (trm_app (trm_app val_add
(trm_var "n")) (val_int 1)) (
trm_app (trm_app val_set (trm_var "p")) (trm_var "m")))).
Alternatively, using notation and coercions, the same program can be written more concisely.
Definition incr' : val :=
<{ fun 'p
let 'n = ! 'p in
let 'm = 'n + 1 in
'p := 'm }>.
Let us check that the two definitions are indeed the same.
Lemma incr_eq_incr' :
incr = incr'.
Proof using. reflexivity. Qed.
Recall from the first chapter the specification of the increment function. Its precondition requires a singleton state of the form p ~~> n. Its postcondition describes a state of the form p ~~> (n+1).
Lemma triple_incr : (p:loc) (n:int),
triple (trm_app incr p)
(p ~~> n)
(fun _p ~~> (n+1)).
We next show a detailed proof for this specification. It exploits:
• the structural reasoning rules,
• the reasoning rules for terms,
• the specification of the primitive functions,
• the xsimpl tactic for simplifying entailments.
Proof using.
(* initialize the proof *)
intros. applys triple_app_fun. { reflexivity. } simpl.
(* reason about let n = .. *)
applys triple_let.
(* reason about !p *)
{ apply triple_get. }
(* name n' the result of !p *)
intros n'. simpl.
(* substitute away the equality n' = n *)
apply triple_hpure. intros →.
(* reason about let m = .. *)
applys triple_let.
(* apply the frame rule to put aside p ~~> n *)
{ applys triple_conseq_frame.
(* reason about n+1 in the empty state *)
{ applys triple_add. }
{ xsimpl. }
{ xsimpl. } }
(* name m' the result of n+1 *)
intros m'. simpl.
(* substitute away the equality m' = m *)
apply triple_hpure. intros →.
(* reason about p := m *)
applys triple_conseq_frame.
{ applys triple_set. }
{ xsimpl. }
{ xsimpl. }
Qed.

### Proof of succ_using_incr

Recall from Basic the function succ_using_incr.
Recall the specification of succ_using_incr.
Lemma triple_succ_using_incr : (n:int),
triple (trm_app succ_using_incr n)
\[]
(fun v\[v = val_int (n+1)]).

#### Exercise: 4 stars, standard, especially useful (triple_succ_using_incr)

Verify the function triple_succ_using_incr. Hint: follow the pattern of triple_incr. Use applys triple_seq for reasoning about a sequence. Use applys triple_val for reasoning about the final return value, namely x.
Proof using. (* FILL IN HERE *) Admitted.

## What's Next

The matter of the next chapters is to introduce additional technology to streamline the proof process, notably by:
• automating the application of the frame rule,
• eliminating the need to manipulate program variables and substitutions during the verification proof.

# More Details

The "more details" section presents several additional reasoning rules.

## Triple for Terms with Same Semantics

A general principle is that if, according to the semantics, the terms t1 and t2 admits the same behaviors, then t1 and t2 satisfy the same triples. The statement can also be reformulated in an asymmetric fashion: if t2 admits no more behaviors than t1, then t2 satisfies all the triples that t1 satisfies. This section explains how to formalize this asymmetric statement, which is more general and more useful in practice than the statement expressed in terms of equivalences.
The judgment eval_like t1 t2 asserts that if t1 terminates and produces results in Q, then so does t2.
Definition eval_like (t1 t2:trm) : Prop :=
s Q, eval s t1 Q eval s t2 Q.

#### Exercise: 1 star, standard, especially useful (eval_like_eta_expansion)

Prove that when the relation eval_like t1 t2 holds, any triple that t2 satisfies any triple that t1 satisfies.
Lemma triple_eval_like : t1 t2 H Q,
eval_like t1 t2
triple t1 H Q
triple t2 H Q.
Proof using. (* FILL IN HERE *) Admitted.
The remaining of this section presents 4 examples applications of this reasoning principle. To begin with, consider the term let x = t in x, represented as trm_let x t x for a variable named x. This term reduces to t, hence the relation eval_like t (trm_let x t x) holds.
Lemma eval_like_eta_reduction : (t:trm) (x:var),
eval_like t (trm_let x t x).
Proof using.
introv R. applys eval_let R.
simpl. rewrite var_eq_spec. case_if.
introv Hv. applys eval_val Hv.
Qed.

#### Exercise: 4 stars, standard, optional (eval_like_eta_expansion)

Prove that the symmetric relation eval_like (trm_let x t x) t also holds.
Lemma eval_like_eta_expansion : (t:trm) (x:var),
eval_like (trm_let x t x) t.
Proof using. (* FILL IN HERE *) Admitted.

#### Exercise: 2 stars, standard, especially useful (eta_same_triples)

Conclude that t and trm_let x t x satisfy exactly the same set of triples. (This result will be exploited in chapter Affine, for proving the equivalence of two versions of the garbage collection rule.)
Lemma eta_same_triples : (t:trm) (x:var) H Q,
triple t H Q triple (trm_let x t x) H Q.
Proof using. (* FILL IN HERE *) Admitted.
Another application of the reasoning rule triple_eval_like is to revisit the proof of triple_app_fun by arguing that trm_app (val_fun x t1) v2 reduces to subst x v2 t1, hence they admit the same behavior.
Lemma triple_app_fun : x v1 v2 t1 H Q,
v1 = val_fun x t1
triple (subst x v2 t1) H Q
triple (trm_app v1 v2) H Q.
Proof using.
introv E M1. applys triple_eval_like M1.
introv R. applys eval_app_fun E R.
Qed.
A even more interesting application is a rule that allows to modify the parenthesis structure of a sequence, from t1; (t2; t3) to (t1;t2); t3. Such a change in the parenthesis structure of a sequence might be helfpul to apply the frame rule around t1;t2, for example.

#### Exercise: 3 stars, standard, optional (triple_trm_seq_assoc)

Prove that the term t1; (t2; t3), which corresponds to the natural parsing of t1; t2; t3, satisfies any triple that the term (t1;t2); t3 satisfies.
Lemma triple_trm_seq_assoc : t1 t2 t3 H Q,
triple (trm_seq (trm_seq t1 t2) t3) H Q
triple (trm_seq t1 (trm_seq t2 t3)) H Q.
Proof using. (* FILL IN HERE *) Admitted.

## The Combined Let-Frame Rule

Module LetFrame.
Recall the Separation Logic let rule.
Parameter triple_let : x t1 t2 Q1 H Q,
triple t1 H Q1
( v1, triple (subst x v1 t2) (Q1 v1) Q)
triple (trm_let x t1 t2) H Q.
At first sight, it seems that, to reason about let x = t1 in t2 in a state described by precondition H, we need to first reason about t1 in that same state. Yet, t1 may well require only a subset of the state H to evaluate, and not all of H.
The "let-frame" rule combines the rule for let-bindings with the frame rule to make it more explicit that the precondition H may be decomposed in the form H1 \* H2, where H1 is the part needed by t1, and H2 denotes the rest of the state. The part of the state covered by H2 remains unmodified during the evaluation of t1, and appears as part of the precondition of t2. The corresponding statement is as follows.

#### Exercise: 3 stars, standard, especially useful (triple_let_frame)

Prove the let-frame rule.
Proof using. (* FILL IN HERE *) Admitted.
End LetFrame.

# Optional Material

## Alternative Specification Style for Pure Preconditions

Module DivSpec.
Recall the specification for division.
Parameter triple_div : n1 n2,
n2 0
triple (val_div n1 n2)
\[]
(fun r\[r = val_int (Z.quot n1 n2)]).
Equivalently, we could place the requirement n2 0 in the precondition:
Parameter triple_div' : n1 n2,
triple (val_div n1 n2)
\[n2 0]
(fun r\[r = val_int (Z.quot n1 n2)]).
The two presentations are equivalent. First, let us prove triple_div' using triple_div.
Lemma triple_div'_from_triple_div : n1 n2,
triple (val_div n1 n2)
\[n2 0]
(fun r\[r = val_int (Z.quot n1 n2)]).
Proof using.
intros. applys triple_hpure'. applys triple_div.
Qed.

#### Exercise: 2 stars, standard, especially useful (triple_div_from_triple_div')

Prove triple_div using triple_div'.
Lemma triple_div_from_triple_div' : n1 n2,
n2 0
triple (val_div n1 n2)
\[]
(fun r\[r = val_int (Z.quot n1 n2)]).
Proof using. (* FILL IN HERE *) Admitted.
Placing pure preconditions outside of the triples makes it slightly more convenient to exploit specifications. For this reason, we adopt the style that precondition only contain the description of heap-allocated data structures.
End DivSpec.

## Alternative Specification Style for Result Values

Module MatchStyle.
Recall the specification for the function ref.
Parameter triple_ref : v,
triple (val_ref v)
\[]
(fun r\ p, \[r = val_loc p] \* p ~~> v).
Its postcondition could be equivalently stated by using, instead of an existential quantifier, a pattern matching construct.
Parameter triple_ref' : v,
triple (val_ref v)
\[]
(fun rmatch r with
| val_loc p ⇒ (p ~~> v)
| _\[False]
end).
However, the pattern-matching presentation is less readable and would be fairly cumbersome to work with in practice.
End MatchStyle.

### Proof of factorec

This section presents a manual proof for a recursive function.
Module ExamplePrograms2.
Export ProgramSyntax.
Import Basic.Facto.
In this section, the corollary triple_hpure' established in the previous chapter, can be useful.
Parameter triple_hpure' : t (P:Prop) Q,
(P triple t \[] Q)
triple t \[P] Q.
Recall the function repeat_incr from chapter Basic.
let rec factorec n =
if n ≤ 1 then 1 else n * factorec (n-1)

#### Exercise: 4 stars, standard, especially useful (triple_factorec)

Verify the function factorec. Hint: exploit triple_app_fix for reasoning about the recursive function. Use triple_hpure', the corollary of triple_hpure. Exploit triple_le and triple_sub and triple_mul to reason about the behavior of the primitive operations involved. Exploit applys triple_if. case_if as C. to reason about the conditional; alternatively, if using triple_if_case, you'll need to use the tactic rew_bool_eq in * to simplify, e.g., the expression isTrue (m 1)) = true.
Lemma triple_factorec : n,
n 0
triple (factorec n)
\[]
(fun r\[r = facto n]).
Proof using. (* FILL IN HERE *) Admitted.

## Historical Notes

[Gordon 1989] presents the first mechanization of Hoare logic in a proof assistant, using the HOL tool. Gordon's pioneering work was followed by numerous formalizations of Hoare logic, targeting various programming languages.
The original presentation of Separation Logic (1999-2001) consists of a set of rules written down on paper. These rules were not formally described in a proof assistant. Nevertheless, mechanized presentation of Separation Logic emerged a few years later.
[Yu, Hamid, and Shao 2003] present the CAP framework for the verification in Coq of assembly-level code. This framework exploits separation logic style specifications, with predicate for lists and list segments involving the separating conjunction operator.
In parallel, [Weber 2004], advised by Nipkow, developed the first mechanization of the rules of Separation Logic for a while language, using the Isabelle/HOL tool. His presentation is quite close from the original, paper presentation.
Numerous mechanized presentations of Separation Logic, targeting various languages (assembly, C, core-Java, ML, etc.) and using various tools (Isabelle/HOL, Coq, PVS, HOL4, HOL). For a detailed list, we refer to the last chapter of the companion notes, linked from the Preface.
(* 2024-01-03 14:19 *)