I am trying to prove the following lemma in Coq:
Lemma not_eq_S2: forall m n, S m <> S n -> m <> n.
It seems easy but I do not find how to finish the proof. Can anybody help me please?
Thank you.
The thing to know is that in Coq, a negation is a function that implies False, so the S m <> S n is really S m = S n -> False. So instead of proving n <> m we can introduce the n = m (we can either unfold not or tell intros explicitly to do it) and get the goal False instead. But with n = m in the context we can rewrite HS: S n <> S m into HS: S n <> S n, which can be handled by auto, or many other tactics such as apply HS. reflexivity. or congruence. etc.
Lemma not_eq_S2: forall m n, S m <> S n -> m <> n.
Proof. intros m n HS HC.
rewrite HC in HS. auto.
Qed.
It's really easy (but the negation makes it a bit confusing).
Lemma not_eq_S2: forall m n, S m <> S n -> m <> n.
Proof.
unfold not. (* |- ... -> False *)
intros m n H C. (* ..., H : S m = S n -> False |- False *)
apply H. (* ... |- S m = S n *)
(* f_equal gets rid of functions applied on both sides of an equality,
this is probably what you didn't find *)
(* basically, f_equal turns a goal [f a = f b] into [a = b] *)
f_equal. (* ..., C : m = n |- m = n *)
exact C.
Qed.
Related
Coq's standard libraries give the Peano natural numbers and addition:
Inductive nat : Set :=
| O : nat
| S : nat -> nat.
Fixpoint add n m :=
match n with
| 0 => m
| S p => S (add p m)
end.
I am curious if I change the fix_definition of addition like
Fixpoint add n m :=
match n with
| 0 => m
| S p => add p (S m)
end.
Is the new addition equivalent to the old one? I tried to prove their equivalence by proving forall n m, add (S n) m = S (add n m) but failed.
In order to proof your helper lemma, you need to be careful what to introduce. If you don't introduce m, you get a more general induction hypothesis as in:
Require Import Nat.
Print add.
Fixpoint my_add n m :=
match n with
| 0 => m
| S p => my_add p (S m)
end.
Lemma my_add_S_r: forall n m, my_add n (S m) = S (my_add n m).
Proof.
(* Note: don't introduce m here - you get a more general induction hypothesis this way *)
intros n.
induction n.
- intros; reflexivity.
- intros; cbn. rewrite IHn. reflexivity.
Qed.
Lemma my_add_equiv: forall n m, add n m = my_add n m.
intros.
induction n.
- reflexivity.
- cbn. rewrite my_add_S_r. rewrite IHn. reflexivity.
Qed.
Yes both additions are equivalent, you can prove it using the lemma plus_n_Sm : forall n m : nat, S (n + m) = n + S m from the standard library (found using Search "+" (S _).) and an adequate induction hypothesis (for instance P(n) := forall m, n + m = add n m).
I'm trying to prove eqb_neq:
Theorem eqb_neq : forall x y : nat,
x =? y = false <-> x <> y.
This is my current proof status:
During the proof I reached a final step where I just need to prove the additional helper theorem:
Theorem eqb_false_helper : forall n m : nat,
n <> m -> S n <> S m.
I've tried multiple strategies but now I'm not even sure it's possible to prove this helper theorem.
I'm not sure how to prove the base case using induction:
What else can I try? Any tips for eqb_neq or the helper theorem?
Thanks
It is actually quite simple for your helper theorem if you just unfold not :
Theorem eqb_false_helper : forall n m : nat,
n <> m -> S n <> S m.
Proof.
unfold not; intros.
apply H; injection H0; intro; assumption.
Qed.
You actually just need to prove that S n = S m -> False, you assume that n = m -> False, thus you can prove that S n = S m -> n = m, which is done injecting hypothesis S n = S m.
I am working on the proof of the following theorem Sn_le_Sm__n_le_m in IndProp.v of Software Foundations (Vol 1: Logical Foundations).
Theorem Sn_le_Sm__n_le_m : ∀n m,
S n ≤ S m → n ≤ m.
Proof.
intros n m HS.
induction HS as [ | m' Hm' IHm'].
- (* le_n *) (* Failed Here *)
- (* le_S *) apply IHSm'.
Admitted.
where, the definition of le (i.e., ≤) is:
Inductive le : nat → nat → Prop :=
| le_n n : le n n
| le_S n m (H : le n m) : le n (S m).
Notation "m ≤ n" := (le m n).
Before induction HS, the context as well as the goal is as follows:
n, m : nat
HS : S n <= S m
______________________________________(1/1)
n <= m
At the point of the first bullet -, the context as well as the goal is:
n, m : nat
______________________________________(1/1)
n <= m
where we have to prove n <= m without any context, which is obviously impossible.
Why does it not generate S n = S m (and then n = m) for the le_n case in induction HS?
The main problem here -I think- is it is impossible to prove the Theorem using induction on HS as there is no way to say something about n with only hypothesis about S n because non of the constructors of le do not change the value of n. But anyway the reason that after first bullet - there is no assumption is because calling induction has the effect of replacing all occurrences of the property argument by the values that correspond to each constructor and it doesn't help in this case since the term that gets replaced S n is not mentioned anywhere. There are some tricks to avoid this. for example you can replace n with pred(S n) as follows.
Theorem Sn_le_Sm__n_le_m : forall n m,
S n <= S m -> n <= m.
Proof.
intros n m HS.
assert(Hn: n=pred (S n)). reflexivity. rewrite Hn.
assert(Hm: m=pred (S m)). reflexivity. rewrite Hm.
induction HS.
- (* le_n *) apply le_n.
- (* le_S *) (* Stucks! *) Abort.
But as I mentioned above it is impossible to go further. Another way is to use inversion which is smarter but in some cases it may not help since induction hypothesis would be necessary. But it worth to know about it.
Theorem Sn_le_Sm__n_le_m : forall n m,
S n <= S m -> n <= m.
Proof.
intros n m HS.
inversion HS.
- (* le_n *) apply le_n.
- (* le_S *) (* Stucks! *) Abort.
Best way to solve the problem is use of remember tactic as follows.
Theorem Sn_le_Sm__n_le_m : forall n m,
S n <= S m -> n <= m.
Proof.
intros n m HS.
remember (S n) as Sn.
remember (S m) as Sm.
induction HS as [ n' | n' m' H IH].
- (* le_n *)
rewrite HeqSn in HeqSm. injection HeqSm as Heq.
rewrite <- Heq. apply le_n.
- (* le_S *) (* Stucks! *) Abort.
According to Software Foundations (Vol 1: Logical Foundations)
The tactic remember e as x causes Coq to (1) replace all occurrences
of the expression e by the variable x, and (2) add an equation x = e
to the context.
Anyway, although it is impossible to prove the fact using induction on HS -imo-, performing an induction on m will solve the case. (Note the use of inversion.)
Theorem Sn_le_Sm__n_le_m : forall n m,
S n <= S m -> n <= m.
Proof.
intros n.
induction m as [|m' IHm'].
- intros H. inversion H as [Hn | n' contra Hn'].
+ apply le_n.
+ inversion contra.
- intros H. inversion H as [HnSm' | n' HSnSm' Heq].
+ apply le_n.
+ apply le_S. apply IHm'. apply HSnSm'.
Qed.
Just more examples of Kamyar's answer.
Well, let's take a look of le induction scheme :
Compute le_ind.
forall (n : nat) (P : nat -> Prop),
P n ->
(forall m : nat, n <= m -> P m -> P (S m)) ->
forall n0 : nat, n <= n0 -> P n0
P is some proposition that holds one natural number, which means in the case of le_n, our preposition n <= m will be reduced to forall n, n <= m. Indeed, it's the same lemma that we want to prove, however unprovable because there is no premise.
An easy to solve this is doing induction where le_ind doesn't do.
For example :
Theorem Sn_le_Sm__n_le_m' : forall m n,
S n <= S m -> n <= m.
elim.
by intros; apply : Gt.gt_S_le .
intros; inversion H0.
by subst.
by subst; apply : le_Sn_le.
Qed.
Notice that we doing induction by m, and using inversion to generates the two possible construction of le ({x = y} + {x < y}). Optionally, you can use le decidability.
Theorem Sn_le_Sm__n_le_m : forall n m,
S n <= S m -> n <= m.
intros.
generalize dependent n.
elim.
auto with arith.
intros.
have : n <= m.
by apply : H; apply : le_Sn_le.
move => H'.
destruct m.
auto with arith.
destruct (le_lt_eq_dec _ _ H').
assumption.
subst.
(* just prove that there is no S m <= m *)
Qed.
For the sake of your time, coq has the tactic dependent induction that easily solves your goal :
Theorem Sn_le_Sm__n_le_m'' : forall n m,
S n <= S m -> n <= m.
intros.
dependent induction H.
auto.
by apply : (le_Sn_le _ _ H).
Qed.
After destructing n in my proof, I am stuck at the following:
1 subgoal
n : nat
X : Type
h : X
t : list X
n' : nat
E : n = S n'
H' : length t = n'
IHl : length t = n -> nth_error t n = None
______________________________________(1/1)
nth_error t n' = None
I want to rewrite using IHl, but that is not possible. How do I compose IHl and H' to make sense and prove this theorem?
I am just trying to elaborate on #Arthur answer.
I was able to reproduce your goal with the following script:
Require Import List.
Lemma toto (n : nat) (X : Type) (l : list nat) : length l = n -> nth_error l n = None.
Proof.
induction l as [ | h t IHl].
case_eq n.
simpl; auto.
simpl; discriminate.
case_eq n.
simpl; discriminate.
intros n' E.
simpl; intros E'; injection E'; clear E'; intros H'.
and I agree that this goal cannot be proved. Now, if you instead start your proof with the following text (the Proof and induction lines have to be replaced), it will be provable (I checked).
Proof.
revert n.
induction l as [ | h t IHl]; intros n.
The difference is that the induction hypothesis now has the following statement.
forall n : nat, length t = n -> nth_error t n = None
What happened? In the first (faulty) variant, you attempt to prove a statement for all lists whose length is equal to a precise n, because n is fixed before you start the proof by induction. In the second (correct) variant, you attempt to prove a statement for all lists l, and this statement accepts any n as long as length l = n.
In the first variant, n is fixed and the equality length l = n restricts l to be among those that have length precisely n. In the second case, l is chosen first, and n is not fixed, but the equality length l = n restricts n to follow the length of l.
This is called "loading the induction" because the statement forall n, length l = n -> nth_error l n = None is stronger (it is loaded) than the statement that you attempt to prove in the first variant (just for one specific n), but surprisingly it is easier to prove.
You cannot, because your induction hypothesis is not general enough. Here is a statement that should be easier to prove:
forall (X : Type) (t : list X), nth_error t (length t) = None
Suppose I want to prove following Theorem:
Theorem succ_neq_zero : forall n m: nat, S n = m -> 0 = m -> False.
This one is trivial since m cannot be both successor and zero, as assumed. However I found it quite tricky to prove it, and I don't know how to make it without an auxiliary lemma:
Lemma succ_neq_zero_lemma : forall n : nat, O = S n -> False.
Proof.
intros.
inversion H.
Qed.
Theorem succ_neq_zero : forall n m: nat, S n = m -> 0 = m -> False.
Proof.
intros.
symmetry in H.
apply (succ_neq_zero_lemma n).
transitivity m.
assumption.
assumption.
Qed.
I am pretty sure there is a better way to prove this. What is the best way to do it?
You just need to substitute for m in the first equation:
Theorem succ_neq_zero : forall n m: nat, S n = m -> 0 = m -> False.
Proof.
intros n m H1 H2; rewrite <- H2 in H1; inversion H1.
Qed.
There's a very easy way to prove it:
Theorem succ_neq_zero : forall n m: nat, S n = m -> 0 = m -> False.
Proof.
congruence.
Qed.
The congruence tactic is a decision procedure for ground equalities on uninterpreted symbols. It's complete for uninterpreted symbols and for constructors, so in cases like this one, it can prove that the equality 0 = m is impossible.
It might be useful to know how congruence works.
To prove that two terms constructed by different constructors are in fact different, just create a function that returns True in one case and False in the other cases, and then use it to prove True = False. I think this is explained in Coq'Art
Example not_congruent: 0 <> 1.
intros C. (* now our goal is 'False' *)
pose (fun m=>match m with 0=>True |S _=>False end) as f.
assert (Contra: f 1 = f 0) by (rewrite C; reflexivity).
now replace False with True by Contra.
Qed.