Proof (not-quite) irrelevance when when destructing coq equality - coq

I have a dependent type which models a finite path in a transition system. The transition system has a function R that yields a proposition saying whether there's an edge between states s and s' with label a. The finite path type is:
Inductive FinPathTail (s : S i) :=
| FPTNil: FinPathTail s
| FPTCons (a : Act i) (s' : S i) : R i s a s' ->
FinPathTail s' -> FinPathTail s.
(The "tail" bit is because this actually models all but the head of a path starting at s).
I've defined a CoInductive type for a possibly infinite PathTail (I'll stick it at the bottom so as to get to the question faster) and I have a function, fpt_to_pt, to transform a FinPathTail into a PathTail. This should "obviously" be injective, so I wanted to prove a lemma of this form:
Lemma fpt_to_pt_inj {s : S i} (fpt fpt' : FinPathTail s)
: (forall s s' : S i, {s = s'} + {s <> s'}) ->
fpt_to_pt fpt = fpt_to_pt fpt' -> fpt = fpt'.
When trying to prove this by induction on fpt, I quickly get to the case where both sides are known to be conses. The goal ends up looking something like:
PTCons s a s' r (fpt_to_pt fpt) = PTCons s a2 s'2 r2 (fpt_to_pt fpt') ->
FPTCons s a s' r fpt = FPTCons s a2 s'2 r2 fpt'
that I'd like to decompose with the injection tactic. The result ends up like this:
existT (fun s'0 : S i => PathTail s'0) s' (fpt_to_pt fpt) =
existT (fun s'0 : S i => PathTail s'0) s'2 (fpt_to_pt fpt') ->
s' = s'2 -> a = a2 -> FPTCons s a s' r fpt = FPTCons s a2 s'2 r2 fpt'
and using the inversion_sigma tactic, I can transform it to:
B : s' = s'2
C : a = a2
A0 : s' = s'2
A1 : eq_rect s' (fun a : S i => PathTail a) (fpt_to_pt fpt) s'2 A0 = fpt_to_pt fpt'
I think I understand why I need decidability for the source domain, in order to use inj_pair2_eq_dec. What I don't understand is: what happened to r and r2? I understand that I don't have proof irrelevance, but doesn't that mean that they must have been equal in order for the conses to be equal? Or am I misunderstanding something fundamental about propositions?
PS: Here's the coinductive definition for PathTail:
CoInductive PathTail (s : S i) :=
| PTNil: PathTail s
| PTCons (a : Act i) (s' : S i) : R i s a s' -> PathTail s' -> PathTail s.

Apparently the injection tactic ignores equalities between proofs by default, but you can override this behavior with the Keep Proof Equalities flag:
Inductive foo : nat -> Prop :=
| Foo (n : nat) : foo n.
Inductive bar :=
| Bar (n : nat) : foo n -> bar.
Lemma test n nn m mm : Bar n nn = Bar m mm -> False.
Proof.
intros H. injection H. (* No equality generated. *)
Abort.
Set Keep Proof Equalities.
Lemma test n nn m mm : Bar n nn = Bar m mm -> False.
Proof.
intros H. injection H. (* Equality generated. *)
Abort.

Related

How to group duplicated hypothesis in Coq?

I have
1 subgoals, subgoal 1
n : nat
b : bool
m : nat
H1: P1
H2: P2
H3: P1
H4: P2
=========
some_goal
after I run the tactic auto_group_duplicates,
it will become
1 subgoals, subgoal 1
n, m : nat
b : bool
H1, H3: P1
H2, H4: P2
=========
some_goal
Is there a tactic like this one?
I don't think that there is a tactic like this. But you can always come up with something using Ltac.
From Coq Require Import Utf8.
Definition mark {A : Type} (x : A) := x.
Ltac bar :=
match goal with
| x : ?A, y : ?A |- _ =>
lazymatch A with
| context [ mark ] => fail
| _ =>
move y after x ;
change A with (mark A) in x
end
end.
Ltac remove_marks :=
repeat match goal with
| x : mark ?A |- _ =>
change (mark A) with A in x
end.
Ltac auto_group_duplicates :=
repeat bar ; remove_marks.
Lemma foo :
∀ (n : nat) (b : bool) (m : nat) (c : bool),
n = m →
b = c →
n = m →
b = c →
n = m →
True.
Proof.
intros n b m c e1 e2 e3 e4 e5.
auto_group_duplicates.
auto_group_duplicates.
Here I have to apply the tactic twice because Ltac unifies A with mark A annoyingly.
You can manually use the move tactic (see https://coq.inria.fr/refman/proof-engine/tactics.html?highlight=top#coq:tacn.move-%E2%80%A6-after-%E2%80%A6) to rearrange hypothesis.
I doubt that there is a general automatic tactic like the one you are looking for, but one could probably cook up a fancy match-goal-based one to automate rearranging hypothesis using this tactic as a basic block, although this is out of my league.

Is there any thing like apply lem in *?

Is there any way to call apply lem in H for every possible H in premises, like rewrite lem in *?
Axiom P Q : nat -> Prop.
Axiom lem : forall (n : nat), P n -> Q n.
Goal P O -> P (S O) -> True.
intros. apply lem in H. apply lem in H0.
I couldn't find anything built in, but it's possible to write such a tactic with Ltac.
First, the special case.
Axiom P Q : nat -> Prop.
Axiom lem : forall (n : nat), P n -> Q n.
Goal P O -> P (S O) -> True.
intros.
repeat match goal with
x : _ |- _ => apply lem in x
end.
Abort.
Now we can generalize this
Ltac apply_in_premises t :=
repeat match goal with
x : _ |- _ => apply t in x
end.
and use it like this:
Goal P O -> P (S O) -> True.
intros.
apply_in_premises lem.
Abort.
Unfortunately, this way of doing it can cause an infinite loop if applying lem produces something else that lem can be applied to.
Axiom P : nat -> Prop.
Axiom lem : forall (n : nat), P n -> P (S n).
Ltac apply_in_premises t :=
repeat match goal with
x : _ |- _ => apply t in x
end.
Goal P O -> P (S O) -> nat -> True.
intros.
apply_in_premises lem. (* infinite loop *)
Abort.
If this is a concern for you, you can use a variant suggested by Yves in the comments. Simply changing apply t in x to apply t in x; revert x will ensure that that hypothesis won't be matched again. However, the end result will have all the hypotheses in the goal, like P -> G, instead of p: P as a premise and G as the goal.
To automatically reintroduce these hypotheses, we can keep track of how many times a hypothesis was reverted, then introduce them again.
Ltac intro_n n :=
match n with
| 0 => idtac
| S ?n' => intro; intro_n n'
end.
Ltac apply_in_premises_n t n :=
match goal with
| x : _ |- _ => apply t in x; revert x;
apply_in_premises_n t (S n)
| _ => intro_n n (* now intro all the premises that were reverted *)
end.
Tactic Notation "apply_in_premises" uconstr(t) := apply_in_premises_n t 0.
Axiom P : nat -> Prop.
Axiom lem : forall (n : nat), P n -> P (S n).
Goal P O -> P (S O) -> nat -> True.
intros.
apply_in_premises lem. (* only applies `lem` once in each of the premises *)
Abort.
Here, the tactic intro_n n applies intro n times.
I haven't tested this in general, but it works well in the case above. It might fail if a hypothesis can't be reverted (for example, if some other hypothesis depends on it). It also may reorder the hypotheses, since when a reverted hypothesis is reintroduced, it's put on the end of the hypothesis list.

Prove that the only zero-length vector is nil

I have a type defined as
Inductive bits : nat -> Set :=
| bitsNil : bits 0
| bitsCons : forall {l}, bool -> bits l -> bits (S l).
and I'm trying to prove:
Lemma emptyIsAlwaysNil : forall {a: bits 0}, a = bitsNil.
After intros, I've tried constructor 1, case a, intuition, to no avail. case a seems like the closest, but it gets an error:
Abstracting over the terms "0" and "a" leads to a term
fun (n : nat) (a0 : bits n) => a0 = bitsNil
which is ill-typed.
Reason is: Illegal application:
The term "#eq" of type "forall A : Type, A -> A -> Prop"
cannot be applied to the terms
"bits n" : "Set"
"a0" : "bits n"
"bitsNil" : "bits 0"
The 3rd term has type "bits 0" which should be coercible to
"bits n".
It sounds like it can't determine whether a bit-vector of an arbitrary length is equal to one of zero-length, because they're different at the type level. Is that correct?
Yes, you're basically correct: specifically, what isn't type checking is Coq's attempt to construct a match on a:bits 0 (which is what case does): the bitsCons case has an ill-typed conclusion.
Here's an axiom-free proof. The key idea is to manually generalize the statement to any n = 0 (I couldn't figure out how to do this with tactics; they all trip up on the dependency). The equality proof then makes the conclusion type check regardless of what n is, and we can dismiss the bitsCons case because we'll have n = S n'. In the more difficult bitsNil case, we make use of eq_rect_eq_dec, which is a consequence of Axiom K but is provable when the type index (nat, in this case) has decidable equality. See the Coq standard library documentation for some other things you can do without axioms with decidable equality.
Require PeanoNat.
Require Import Eqdep_dec.
Import EqNotations.
Inductive bits : nat -> Set :=
| bitsNil : bits 0
| bitsCons : forall {l}, bool -> bits l -> bits (S l).
Lemma emptyIsAlwaysNil_general :
forall n (H: n = 0) {a: bits n},
rew [bits] H in a = bitsNil.
Proof.
intros.
induction a; simpl.
(* bitsNil *)
rewrite <- eq_rect_eq_dec; auto.
apply PeanoNat.Nat.eq_dec.
(* bitsCons - derive a contradiction *)
exfalso; discriminate H.
Qed.
Lemma emptyIsAlwaysNil : forall {a: bits 0},
a = bitsNil.
Proof.
intros.
change a with (rew [bits] eq_refl in a).
apply emptyIsAlwaysNil_general.
Qed.
You don't need the rew H in x notation from EqNotations (it just wraps eq_rect, the equality recursion principle), but I find it makes things much more readable.
However, you can prove this theorem more simply if you're willing to use an axiom, specifically JMeq_eq (see CPDT's equality chapter for more details), since then you can use dependent induction or dependent destruction:
Require Import Program.Equality.
Inductive bits : nat -> Set :=
| bitsNil : bits 0
| bitsCons : forall {l}, bool -> bits l -> bits (S l).
Lemma emptyIsAlwaysNil :
forall {a: bits 0}, a = bitsNil.
Proof.
intros.
dependent destruction a; reflexivity.
Qed.
Print Assumptions emptyIsAlwaysNil.
(* Axioms:
JMeq_eq : forall (A : Type) (x y : A), x ~= y -> x = y *)
Here is a simple proof (borrowed from this Coq Club thread):
Definition emptyIsAlwaysNil {a: bits 0} : a = bitsNil :=
match a with bitsNil => eq_refl end.
Opaque emptyIsAlwaysNil.
Here is what Coq builds under the hood:
Print emptyIsAlwaysNil.
emptyIsAlwaysNil =
fun a : bits 0 =>
match
a as a0 in (bits n)
return
(match n as x return (bits x -> Type) with
| 0 => fun a1 : bits 0 => a1 = bitsNil
| S n0 => fun _ : bits (S n0) => IDProp
end a0)
with
| bitsNil => eq_refl
| bitsCons _ _ => idProp
end
: forall a : bits 0, a = bitsNil

coq induction with passing in equality

I have a list with a known value and want to induct on it, keeping track of what the original list was, and referring to it by element. That is, I need to refer to it by l[i] with varying i instead of just having (a :: l).
I tried to make an induction principle to allow me to do that. Here is a program with all of the unnecessary Theorems replaced with Admitted, using a simplified example. The objective is to prove allLE_countDown using countDown_nth, and have list_nth_rect in a convenient form. (The theorem is easy to prove directly without any of those.)
Require Import Arith.
Require Import List.
Definition countDown1 := fix f a i := match i with
| 0 => nil
| S i0 => (a + i0) :: f a i0
end.
(* countDown from a number to another, excluding greatest. *)
Definition countDown a b := countDown1 b (a - b).
Theorem countDown_nth a b i d (boundi : i < length (countDown a b))
: nth i (countDown a b) d = a - i - 1.
Admitted.
Definition allLE := fix f l m := match l with
| nil => true
| a :: l0 => if Nat.leb a m then f l0 m else false
end.
Definition drop {A} := fix f (l : list A) n := match n with
| 0 => l
| S a => match l with
| nil => nil
| _ :: l2 => f l2 a
end
end.
Theorem list_nth_rect_aux {A : Type} (P : list A -> list A -> nat -> Type)
(Pnil : forall l, P l nil (length l))
(Pcons : forall i s l d (boundi : i < length l), P l s (S i) -> P l ((nth i l d) :: s) i)
l s i (size : length l = i + length s) (sub : s = drop l i) : P l s i.
Admitted.
Theorem list_nth_rect {A : Type} (P : list A -> list A -> nat -> Type)
(Pnil : forall l, P l nil (length l))
(Pcons : forall i s l d (boundi : i < length l), P l s (S i) -> P l ((nth i l d) :: s) i)
l s (leqs : l = s): P l s 0.
Admitted.
Theorem allLE_countDown a b : allLE (countDown a b) a = true.
remember (countDown a b) as l.
refine (list_nth_rect (fun l s _ => l = countDown a b -> allLE s a = true) _ _ l l eq_refl Heql);
intros; subst; [ apply eq_refl | ].
rewrite countDown_nth; [ | apply boundi ].
pose proof (Nat.le_sub_l a (i + 1)).
rewrite Nat.sub_add_distr in H0.
apply leb_correct in H0.
simpl; rewrite H0; clear H0.
apply (H eq_refl).
Qed.
So, I have list_nth_rect and was able to use it with refine to prove the theorem by referring to the nth element, as desired. However, I had to construct the Proposition P myself. Normally, you'd like to use induction.
This requires distinguishing which elements are the original list l vs. the sublist s that is inducted on. So, I can use remember.
Theorem allLE_countDown a b : allLE (countDown a b) a = true.
remember (countDown a b) as s.
remember s as l.
rewrite Heql.
This puts me at
a, b : nat
s, l : list nat
Heql : l = s
Heqs : l = countDown a b
============================
allLE s a = true
However, I can't seem to pass the equality as I just did above. When I try
induction l, s, Heql using list_nth_rect.
I get the error
Error: Abstracting over the terms "l", "s" and "0" leads to a term
fun (l0 : list ?X133#{__:=a; __:=b; __:=s; __:=l; __:=Heql; __:=Heqs})
(s0 : list ?X133#{__:=a; __:=b; __:=s; __:=l0; __:=Heql; __:=Heqs})
(_ : nat) =>
(fun (l1 l2 : list nat) (_ : l1 = l2) =>
l1 = countDown a b -> allLE l2 a = true) l0 s0 Heql
which is ill-typed.
Reason is: Illegal application:
The term
"fun (l l0 : list nat) (_ : l = l0) =>
l = countDown a b -> allLE l0 a = true" of type
"forall l l0 : list nat, l = l0 -> Prop"
cannot be applied to the terms
"l0" : "list nat"
"s0" : "list nat"
"Heql" : "l = s"
The 3rd term has type "l = s" which should be coercible to
"l0 = s0".
So, how can I change the induction principle
such that it works with the induction tactic?
It looks like it's getting confused between
the outer variables and the ones inside the
function. But, I don't have a way to talk
about the inner variables that aren't in scope.
It's very strange, since invoking it with
refine works without issues.
I know for match, there's as clauses, but
I can't figure out how to apply that here.
Or, is there a way to make list_nth_rect use
P l l 0 and still indicate which variables correspond to l and s?
First, you can prove this result much more easily by reusing more basic ones. Here's a version based on definitions of the ssreflect library:
From mathcomp
Require Import ssreflect ssrfun ssrbool ssrnat eqtype seq.
Definition countDown n m := rev (iota m (n - m)).
Lemma allLE_countDown n m : all (fun k => k <= n) (countDown n m).
Proof.
rewrite /countDown all_rev; apply/allP=> k; rewrite mem_iota.
have [mn|/ltnW] := leqP m n.
by rewrite subnKC //; case/andP => _; apply/leqW.
by rewrite -subn_eq0 => /eqP ->; rewrite addn0 ltnNge andbN.
Qed.
Here, iota n m is the list of m elements that counts starting from n, and all is a generic version of your allLE. Similar functions and results exist in the standard library.
Back to your original question, it is true that sometimes we need to induct on a list while remembering the entire list we started with. I don't know if there is a way to get what you want with the standard induction tactic; I didn't even know that it had a multi-argument variant. When I want to prove P l using this strategy, I usually proceed as follows:
Find a predicate Q : nat -> Prop such that Q (length l) implies P l. Typically, Q n will have the form n <= length l -> R (take n l) (drop n l), where R : list A -> list A -> Prop.
Prove Q n for all n by induction.
I do not know if this answers your question, but induction seems to accept with clauses. Thus, you can write the following.
Theorem allLE_countDown a b : allLE (countDown a b) a = true.
remember (countDown a b) as s.
remember s as l.
rewrite Heql.
induction l, s, Heql using list_nth_rect
with (P:=fun l s _ => l = countDown a b -> allLE s a = true).
But the benefit is quite limited w.r.t. the refine version, since you need to specify manually the predicate.
Now, here is how I would have proved such a result using objects from the standard library.
Require Import List. Import ListNotations.
Require Import Omega.
Definition countDown1 := fix f a i := match i with
| 0 => nil
| S i0 => (a + i0) :: f a i0
end.
(* countDown from a number to another, excluding greatest. *)
Definition countDown a b := countDown1 b (a - b).
Theorem countDown1_nth a i k d (boundi : k < i) :
nth k (countDown1 a i) d = a + i -k - 1.
Proof.
revert k boundi.
induction i; intros.
- inversion boundi.
- simpl. destruct k.
+ omega.
+ rewrite IHi; omega.
Qed.
Lemma countDown1_length a i : length (countDown1 a i) = i.
Proof.
induction i.
- reflexivity.
- simpl. rewrite IHi. reflexivity.
Qed.
Theorem countDown_nth a b i d (boundi : i < length (countDown a b))
: nth i (countDown a b) d = a - i - 1.
Proof.
unfold countDown in *.
rewrite countDown1_length in boundi.
rewrite countDown1_nth.
replace (b+(a-b)) with a by omega. reflexivity. assumption.
Qed.
Theorem allLE_countDown a b : Forall (ge a) (countDown a b).
Proof.
apply Forall_forall. intros.
apply In_nth with (d:=0) in H.
destruct H as (n & H & H0).
rewrite countDown_nth in H0 by assumption. omega.
Qed.
EDIT:
You can state an helper lemma to make an even more concise proof.
Lemma Forall_nth : forall {A} (P:A->Prop) l,
(forall d i, i < length l -> P (nth i l d)) ->
Forall P l.
Proof.
intros. apply Forall_forall.
intros. apply In_nth with (d:=x) in H0.
destruct H0 as (n & H0 & H1).
rewrite <- H1. apply H. assumption.
Qed.
Theorem allLE_countDown a b : Forall (ge a) (countDown a b).
Proof.
apply Forall_nth.
intros. rewrite countDown_nth. omega. assumption.
Qed.
The issue is that, for better or for worse, induction seems to assume that its arguments are independent. The solution, then, is to let induction automatically infer l and s from Heql:
Theorem list_nth_rect {A : Type} {l s : list A} (P : list A -> list A -> nat -> Type)
(Pnil : P l nil (length l))
(Pcons : forall i s d (boundi : i < length l), P l s (S i) -> P l ((nth i l d) :: s) i)
(leqs : l = s): P l s 0.
Admitted.
Theorem allLE_countDown a b : allLE (countDown a b) a = true.
remember (countDown a b) as s.
remember s as l.
rewrite Heql.
induction Heql using list_nth_rect;
intros; subst; [ apply eq_refl | ].
rewrite countDown_nth; [ | apply boundi ].
pose proof (Nat.le_sub_l a (i + 1)).
rewrite Nat.sub_add_distr in H.
apply leb_correct in H.
simpl; rewrite H; clear H.
assumption.
Qed.
I had to change around the type of list_nth_rect a bit; I hope I haven't made it false.

Dependent pattern matching in coq

The following code (which is of course not a complete proof) tries to do pattern matching on a dependent product:
Record fail : Set :=
mkFail {
i : nat ;
f : forall x, x < i -> nat
}.
Definition failomat : forall (m : nat) (f : forall x, x < m -> nat), nat.
Proof.
intros.
apply 0.
Qed.
Function fail_hard_omat fl : nat := failomat (i fl) (f fl).
Definition failhard fl : fail_hard_omat fl = 0.
refine ((fun fl =>
match fl with
| mkFail 0 _ => _
| mkFail (S n) _ => _
end) fl).
The error I get when trying to execute this is
Toplevel input, characters 0-125:
Error: Illegal application (Type Error):
The term "mkFail" of type
"forall i : nat, (forall x : nat, x < i -> nat) -> fail"
cannot be applied to the terms
"i" : "nat"
"f0" : "forall x : nat, x < i0 -> nat"
The 2nd term has type "forall x : nat, x < i0 -> nat"
which should be coercible to "forall x : nat, x < i -> nat".
It seems that the substitution somehow does not reach the inner type parameters.
After playing with the Program command I managed to build a refine that might suites you, but I don't understand everything I did. The main idea is to help Coq with the substitution by introducing intermediate equalities that will serve as brige within the substitution
refine ((fun fl =>
match fl as fl0 return (fl0 = fl -> fail_hard_omat fl0 = 0) with
| mkFail n bar =>
match n as n0 return (forall foo: (forall x:nat, x < n0 -> nat),
mkFail n0 foo = fl -> fail_hard_omat (mkFail n0 foo) = 0) with
| O => _
| S p => _
end bar
end (eq_refl fl) ) fl).
Anyway, I don't know what your purpose here is, but I advise never write dependent match "by hand" and rely on Coq's tactics. In your case, if you define your Definition failomat with Defined. instead of Qed, you will be able to unfold it and you won't need dependent matching.
Hope it helps,
V.
Note: both occurences of bar can be replaced by an underscore.
Another, slightly less involved, alternative is to use nat and fail's induction combinators.
Print nat_rect.
Print fail_rect.
Definition failhard : forall fl, fail_hard_omat fl = 0.
Proof.
refine (fail_rect _ _). (* Performs induction (projection) on fl. *)
refine (nat_rect _ _ _). (* Performs induction on fl's first component. *)
Show Proof.