# LogicLogic in Coq

We have seen many examples of factual claims (
Before diving into details, let's talk a bit about the status of
mathematical statements in Coq. Recall that Coq is a

*propositions*) and ways of presenting evidence of their truth (*proofs*). In particular, we have worked extensively with*equality propositions*(e_{1}= e_{2}), implications (P → Q), and quantified propositions (∀ x, P). In this chapter, we will see how Coq can be used to carry out other familiar forms of logical reasoning.*typed*language, which means that every sensible expression in its world has an associated type. Logical claims are no exception: any statement we might try to prove in Coq has a type, namely Prop, the type of*propositions*. We can see this with the Check command:
Note that
Simply

*all*syntactically well-formed propositions have type Prop in Coq, regardless of whether they are true.*being*a proposition is one thing; being*provable*is something else!
Indeed, propositions not only have types: they are
So far, we've seen one primary place that propositions can appear:
in Theorem (and Lemma and Example) declarations.

*first-class*entities that can be manipulated in all the same ways as any of the other things in Coq's world.
But propositions can be used in many other ways. For example, we
can give a name to a proposition using a Definition, just as we
have given names to other kinds of expressions.

We can later use this name in any situation where a proposition is
expected -- for example, as the claim in a Theorem declaration.

We can also write
For instance, the following function takes a number
and returns a proposition asserting that this number is equal to
three:

*parameterized*propositions -- that is, functions that take arguments of some type and return a proposition.
In Coq, functions that return propositions are said to define

*properties*of their arguments. For instance, here's a (polymorphic) property defining the familiar notion of an*injective function*.Definition injective {A B} (f : A → B) :=

∀ x y : A, f x = f y → x = y.

Lemma succ_inj : injective S.

Proof.

intros n m H. injection H as H

_{1}. apply H

_{1}.

Qed.

The equality operator = is also a function that returns a
Prop.
The expression n = m is syntactic sugar for eq n m (defined in
Coq's standard library using the Notation mechanism). Because
eq can be used with elements of any type, it is also
polymorphic:

(Notice that we wrote @eq instead of eq: The type
argument A to eq is declared as implicit, and we need to turn
off the inference of this implicit argument to see the full type
of eq.)

# Logical Connectives

## Conjunction

*conjunction*, or

*logical and*, of propositions A and B is written A ∧ B, representing the claim that both A and B are true.

To prove a conjunction, use the split tactic. It will generate
two subgoals, one for each part of the statement:

Proof.

split.

- (* 3 + 4 = 7 *) reflexivity.

- (* 2 * 2 = 4 *) reflexivity.

Qed.

For any propositions A and B, if we assume that A is true
and that B is true, we can conclude that A ∧ B is also
true.

Lemma and_intro : ∀ A B : Prop, A → B → A ∧ B.

Proof.

intros A B HA HB. split.

- apply HA.

- apply HB.

Qed.

Since applying a theorem with hypotheses to some goal has the
effect of generating as many subgoals as there are hypotheses for
that theorem, we can apply and_intro to achieve the same effect
as split.

Example and_example' : 3 + 4 = 7 ∧ 2 × 2 = 4.

Proof.

apply and_intro.

- (* 3 + 4 = 7 *) reflexivity.

- (* 2 + 2 = 4 *) reflexivity.

Qed.

Example and_exercise :

∀ n m : nat, n + m = 0 → n = 0 ∧ m = 0.

Proof.

(* FILL IN HERE *) Admitted.

☐

∀ n m : nat, n + m = 0 → n = 0 ∧ m = 0.

Proof.

(* FILL IN HERE *) Admitted.

☐

*use*a conjunctive hypothesis to help prove something else -- we employ the destruct tactic.

Lemma and_example2 :

∀ n m : nat, n = 0 ∧ m = 0 → n + m = 0.

Proof.

(* WORKED IN CLASS *)

intros n m H.

destruct H as [Hn Hm].

rewrite Hn. rewrite Hm.

reflexivity.

Qed.

As usual, we can also destruct H right when we introduce it,
instead of introducing and then destructing it:

Lemma and_example2' :

∀ n m : nat, n = 0 ∧ m = 0 → n + m = 0.

Proof.

intros n m [Hn Hm].

rewrite Hn. rewrite Hm.

reflexivity.

Qed.

You may wonder why we bothered packing the two hypotheses n = 0
and m = 0 into a single conjunction, since we could have also
stated the theorem with two separate premises:

Lemma and_example2'' :

∀ n m : nat, n = 0 → m = 0 → n + m = 0.

Proof.

intros n m Hn Hm.

rewrite Hn. rewrite Hm.

reflexivity.

Qed.

For this specific theorem, both formulations are fine. But
it's important to understand how to work with conjunctive
hypotheses because conjunctions often arise from intermediate
steps in proofs, especially in larger developments. Here's a
simple example:

Lemma and_example3 :

∀ n m : nat, n + m = 0 → n × m = 0.

Proof.

(* WORKED IN CLASS *)

intros n m H.

apply and_exercise in H.

destruct H as [Hn Hm].

rewrite Hn. reflexivity.

Qed.

Another common situation with conjunctions is that we know
A ∧ B but in some context we need just A or just B.
In such cases we can do a destruct (possibly as part of
an intros) and use an underscore pattern _ to indicate
that the unneeded conjunct should just be thrown away.

Lemma proj1 : ∀ P Q : Prop,

P ∧ Q → P.

Proof.

intros P Q HPQ.

destruct HPQ as [HP _].

apply HP. Qed.

Theorem and_commut : ∀ P Q : Prop,

P ∧ Q → Q ∧ P.

Proof.

intros P Q [HP HQ].

split.

- (* left *) apply HQ.

- (* right *) apply HP. Qed.

#### Exercise: 2 stars, standard (and_assoc)

(In the following proof of associativity, notice how the*nested*intros pattern breaks the hypothesis H : P ∧ (Q ∧ R) down into HP : P, HQ : Q, and HR : R. Finish the proof from there.)

Theorem and_assoc : ∀ P Q R : Prop,

P ∧ (Q ∧ R) → (P ∧ Q) ∧ R.

Proof.

intros P Q R [HP [HQ HR]].

(* FILL IN HERE *) Admitted.

☐

## Disjunction

*disjunction*, or

*logical or*, of two propositions: A ∨ B is true when either A or B is. (This infix notation stands for or A B, where or : Prop → Prop → Prop.)

Lemma eq_mult_0 :

∀ n m : nat, n = 0 ∨ m = 0 → n × m = 0.

Proof.

(* This pattern implicitly does case analysis on

n = 0 ∨ m = 0 *)

intros n m [Hn | Hm].

- (* Here, n = 0 *)

rewrite Hn. reflexivity.

- (* Here, m = 0 *)

rewrite Hm. rewrite <- mult_n_O.

reflexivity.

Qed.

Conversely, to show that a disjunction holds, it suffices to show
that one of its sides holds. This is done via two tactics, left
and right. As their names imply, the first one requires proving
the left side of the disjunction, while the second requires
proving its right side. Here is a trivial use...

... and here is a slightly more interesting example requiring both
left and right:

Lemma zero_or_succ :

∀ n : nat, n = 0 ∨ n = S (pred n).

Proof.

(* WORKED IN CLASS *)

intros [|n'].

- left. reflexivity.

- right. reflexivity.

Qed.

## Falsehood and Negation

So far, we have mostly been concerned with proving that certain things are*true*-- addition is commutative, appending lists is associative, etc. Of course, we may also be interested in negative results, demonstrating that some given proposition is

*not*true. Such statements are expressed with the logical negation operator ¬.

*principle of explosion*from the Tactics chapter, which asserts that, if we assume a contradiction, then any other proposition can be derived. Following this intuition, we could define ¬ P ("not P") as ∀ Q, P → Q.

Module MyNot.

Definition not (P:Prop) := P → False.

Notation "~ x" := (not x) : type_scope.

Check not : Prop → Prop.

End MyNot.

Since False is a contradictory proposition, the principle of
explosion also applies to it. If we get False into the proof
context, we can use destruct on it to complete any goal:

Theorem ex_falso_quodlibet : ∀ (P:Prop),

False → P.

Proof.

(* WORKED IN CLASS *)

intros P contra.

destruct contra. Qed.

The Latin

*ex falso quodlibet*means, literally, "from falsehood follows whatever you like"; this is another common name for the principle of explosion.#### Exercise: 2 stars, standard, optional (not_implies_our_not)

Show that Coq's definition of negation implies the intuitive one mentioned above:Fact not_implies_our_not : ∀ (P:Prop),

¬ P → (∀ (Q:Prop), P → Q).

Proof.

(* FILL IN HERE *) Admitted.

☐

Notation "x <> y" := (~(x = y)).

The proposition 0 ≠ 1 is exactly the same as
~(0 = 1), that is not (0 = 1), which unfolds to
(0 = 1) → False. (We use unfold not explicitly here
to illustrate that point, but generally it can be omitted.)

To prove an inequality, we may assume the opposite
equality...

intros contra.

... and deduce a contradiction from it. Here, the
equality O = S O contradicts the disjointness of
constructors O and S, so discriminate takes care
of it.

discriminate contra.

Qed.

Qed.

It takes a little practice to get used to working with negation in
Coq. Even though you can see perfectly well why a statement
involving negation is true, it can be a little tricky at first to
make Coq understand it! Here are proofs of a few familiar facts
to get you warmed up.

Theorem not_False :

¬ False.

Proof.

unfold not. intros H. destruct H. Qed.

Theorem contradiction_implies_anything : ∀ P Q : Prop,

(P ∧ ¬P) → Q.

Proof.

(* WORKED IN CLASS *)

intros P Q [HP HNA]. unfold not in HNA.

apply HNA in HP. destruct HP. Qed.

Theorem double_neg : ∀ P : Prop,

P → ~~P.

Proof.

(* WORKED IN CLASS *)

intros P H. unfold not. intros G. apply G. apply H. Qed.

#### Exercise: 2 stars, advanced (double_neg_inf)

Write an informal proof of double_neg:*Theorem*: P implies ~~P, for any proposition P.

(* FILL IN HERE *)

(* Do not modify the following line: *)

Definition manual_grade_for_double_neg_inf : option (nat×string) := None.

☐

#### Exercise: 1 star, advanced (informal_not_PNP)

Write an informal proof (in English) of the proposition ∀ P : Prop, ~(P ∧ ¬P).(* FILL IN HERE *)

(* Do not modify the following line: *)

Definition manual_grade_for_informal_not_PNP : option (nat×string) := None.

☐

Theorem not_true_is_false : ∀ b : bool,

b ≠ true → b = false.

Proof.

intros b H.

destruct b eqn:HE.

- (* b = true *)

unfold not in H.

apply ex_falso_quodlibet.

apply H. reflexivity.

- (* b = false *)

reflexivity.

Qed.

Since reasoning with ex_falso_quodlibet is quite common, Coq
provides a built-in tactic, exfalso, for applying it.

Theorem not_true_is_false' : ∀ b : bool,

b ≠ true → b = false.

Proof.

intros [] H. (* note implicit destruct b here *)

- (* b = true *)

unfold not in H.

exfalso. (* <=== *)

apply H. reflexivity.

- (* b = false *) reflexivity.

Qed.

## Truth

Unlike False, which is used extensively, True is used quite
rarely, since it is trivial (and therefore uninteresting) to prove
as a goal, and it carries no useful information as a hypothesis. But it can be quite useful when defining complex Props using
conditionals or as a parameter to higher-order Props. We will
see examples later on.
The handy "if and only if" connective, which asserts that two
propositions have the same truth value, is simply the conjunction
of two implications.

## Logical Equivalence

Module MyIff.

Definition iff (P Q : Prop) := (P → Q) ∧ (Q → P).

Notation "P <-> Q" := (iff P Q)

(at level 95, no associativity)

: type_scope.

End MyIff.

Theorem iff_sym : ∀ P Q : Prop,

(P ↔ Q) → (Q ↔ P).

Proof.

(* WORKED IN CLASS *)

intros P Q [HAB HBA].

split.

- (* -> *) apply HBA.

- (* <- *) apply HAB. Qed.

Lemma not_true_iff_false : ∀ b,

b ≠ true ↔ b = false.

Proof.

(* WORKED IN CLASS *)

intros b. split.

- (* -> *) apply not_true_is_false.

- (* <- *)

intros H. rewrite H. intros H'. discriminate H'.

Qed.

Theorem or_distributes_over_and : ∀ P Q R : Prop,

P ∨ (Q ∧ R) ↔ (P ∨ Q) ∧ (P ∨ R).

Proof.

(* FILL IN HERE *) Admitted.

☐

P ∨ (Q ∧ R) ↔ (P ∨ Q) ∧ (P ∨ R).

Proof.

(* FILL IN HERE *) Admitted.

☐

## Setoids and Logical Equivalence

A "setoid" is a set equipped with an equivalence relation,
that is, a relation that is reflexive, symmetric, and transitive.
When two elements of a set are equivalent according to the
relation, rewrite can be used to replace one element with the
other. We've seen that already with the equality relation = in
Coq: when x = y, we can use rewrite to replace x with y,
or vice-versa.
Similarly, the logical equivalence relation ↔ is reflexive,
symmetric, and transitive, so we can use it to replace one part of
a proposition with another: if P ↔ Q, then we can use
rewrite to replace P with Q, or vice-versa.
Here is a simple example demonstrating how these tactics work with
iff. First, let's prove a couple of basic iff equivalences.

Lemma mult_0 : ∀ n m, n × m = 0 ↔ n = 0 ∨ m = 0.

Theorem or_assoc :

∀ P Q R : Prop, P ∨ (Q ∨ R) ↔ (P ∨ Q) ∨ R.

Proof.

intros P Q R. split.

- intros [H | [H | H]].

+ left. left. apply H.

+ left. right. apply H.

+ right. apply H.

- intros [[H | H] | H].

+ left. apply H.

+ right. left. apply H.

+ right. right. apply H.

Qed.

intros P Q R. split.

- intros [H | [H | H]].

+ left. left. apply H.

+ left. right. apply H.

+ right. apply H.

- intros [[H | H] | H].

+ left. apply H.

+ right. left. apply H.

+ right. right. apply H.

Qed.

We can now use these facts with rewrite and reflexivity
to give smooth proofs of statements involving equivalences. For
example, here is a ternary version of the previous mult_0
result:

Lemma mult_0_3 :

∀ n m p, n × m × p = 0 ↔ n = 0 ∨ m = 0 ∨ p = 0.

Proof.

intros n m p.

rewrite mult_0. rewrite mult_0. rewrite or_assoc.

reflexivity.

Qed.

The apply tactic can also be used with ↔. When given an
equivalence as its argument, apply tries to guess which direction of
the equivalence will be useful.

Lemma apply_iff_example :

∀ n m : nat, n × m = 0 → n = 0 ∨ m = 0.

Proof.

intros n m H. apply mult_0. apply H.

Qed.

## Existential Quantification

*existential quantification*. To say that there is some x of type T such that some property P holds of x, we write ∃ x : T, P. As with ∀, the type annotation : T can be omitted if Coq is able to infer from the context what the type of x should be.

*witness*of the existential. This is done in two steps: First, we explicitly tell Coq which witness t we have in mind by invoking the tactic ∃ t. Then we prove that P holds after all occurrences of x are replaced by t.

Definition even x := ∃ n : nat, x = double n.

Lemma four_is_even : even 4.

Proof.

unfold even. ∃ 2. reflexivity.

Qed.

Conversely, if we have an existential hypothesis ∃ x, P in
the context, we can destruct it to obtain a witness x and a
hypothesis stating that P holds of x.

Theorem exists_example_2 : ∀ n,

(∃ m, n = 4 + m) →

(∃ o, n = 2 + o).

Proof.

(* WORKED IN CLASS *)

intros n [m Hm]. (* note implicit destruct here *)

∃ (2 + m).

apply Hm. Qed.

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

Prove that "P holds for all x" implies "there is no x for which P does not hold." (Hint: destruct H as [x E] works on existential assumptions!)Theorem dist_not_exists : ∀ (X:Type) (P : X → Prop),

(∀ x, P x) → ¬ (∃ x, ¬ P x).

Proof.

(* FILL IN HERE *) Admitted.

☐

#### Exercise: 2 stars, standard (dist_exists_or)

Prove that existential quantification distributes over disjunction.Theorem dist_exists_or : ∀ (X:Type) (P Q : X → Prop),

(∃ x, P x ∨ Q x) ↔ (∃ x, P x) ∨ (∃ x, Q x).

Proof.

(* FILL IN HERE *) Admitted.

☐

# Programming with Propositions

- If l is the empty list, then x cannot occur in it, so the property "x appears in l" is simply false.

- Otherwise, l has the form x' :: l'. In this case, x occurs in l if either it is equal to x' or it occurs in l'.

Fixpoint In {A : Type} (x : A) (l : list A) : Prop :=

match l with

| [] ⇒ False

| x' :: l' ⇒ x' = x ∨ In x l'

end.

When In is applied to a concrete list, it expands into a
concrete sequence of nested disjunctions.

Example In_example_1 : In 4 [1; 2; 3; 4; 5].

Proof.

(* WORKED IN CLASS *)

simpl. right. right. right. left. reflexivity.

Qed.

Example In_example_2 :

∀ n, In n [2; 4] →

∃ n', n = 2 × n'.

Proof.

(* WORKED IN CLASS *)

simpl.

intros n [H | [H | []]].

- ∃ 1. rewrite <- H. reflexivity.

- ∃ 2. rewrite <- H. reflexivity.

Qed.

(Notice the use of the empty pattern to discharge the last case
We can also prove more generic, higher-level lemmas about In.
Note, in the first, how In starts out applied to a variable and
only gets expanded when we do case analysis on this variable:

*en passant*.)Theorem In_map :

∀ (A B : Type) (f : A → B) (l : list A) (x : A),

In x l →

In (f x) (map f l).

Proof.

intros A B f l x.

induction l as [|x' l' IHl'].

- (* l = nil, contradiction *)

simpl. intros [].

- (* l = x' :: l' *)

simpl. intros [H | H].

+ rewrite H. left. reflexivity.

+ right. apply IHl'. apply H.

Qed.

This way of defining propositions recursively, though convenient
in some cases, also has some drawbacks. In particular, it is
subject to Coq's usual restrictions regarding the definition of
recursive functions, e.g., the requirement that they be "obviously
terminating." In the next chapter, we will see how to define
propositions

*inductively*, a different technique with its own set of strengths and limitations.#### Exercise: 3 stars, standard (In_map_iff)

Theorem In_map_iff :

∀ (A B : Type) (f : A → B) (l : list A) (y : B),

In y (map f l) ↔

∃ x, f x = y ∧ In x l.

Proof.

intros A B f l y. split.

(* FILL IN HERE *) Admitted.

☐

∀ (A B : Type) (f : A → B) (l : list A) (y : B),

In y (map f l) ↔

∃ x, f x = y ∧ In x l.

Proof.

intros A B f l y. split.

(* FILL IN HERE *) Admitted.

☐

Theorem In_app_iff : ∀ A l l' (a:A),

In a (l++l') ↔ In a l ∨ In a l'.

Proof.

intros A l. induction l as [|a' l' IH].

(* FILL IN HERE *) Admitted.

☐

In a (l++l') ↔ In a l ∨ In a l'.

Proof.

intros A l. induction l as [|a' l' IH].

(* FILL IN HERE *) Admitted.

☐

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

Recall that functions returning propositions can be seen as*properties*of their arguments. For instance, if P has type nat → Prop, then P n states that property P holds of n.

*not*just restate the left-hand side of All_In.)

Fixpoint All {T : Type} (P : T → Prop) (l : list T) : Prop

(* REPLACE THIS LINE WITH ":= _your_definition_ ." *). Admitted.

Theorem All_In :

∀ T (P : T → Prop) (l : list T),

(∀ x, In x l → P x) ↔

All P l.

Proof.

(* FILL IN HERE *) Admitted.

☐

#### Exercise: 2 stars, standard, optional (combine_odd_even)

Complete the definition of the combine_odd_even function below. It takes as arguments two properties of numbers, Podd and Peven, and it should return a property P such that P n is equivalent to Podd n when n is odd and equivalent to Peven n otherwise.Definition combine_odd_even (Podd Peven : nat → Prop) : nat → Prop

(* REPLACE THIS LINE WITH ":= _your_definition_ ." *). Admitted.

To test your definition, prove the following facts:

Theorem combine_odd_even_intro :

∀ (Podd Peven : nat → Prop) (n : nat),

(oddb n = true → Podd n) →

(oddb n = false → Peven n) →

combine_odd_even Podd Peven n.

Proof.

(* FILL IN HERE *) Admitted.

Theorem combine_odd_even_elim_odd :

∀ (Podd Peven : nat → Prop) (n : nat),

combine_odd_even Podd Peven n →

oddb n = true →

Podd n.

Proof.

(* FILL IN HERE *) Admitted.

Theorem combine_odd_even_elim_even :

∀ (Podd Peven : nat → Prop) (n : nat),

combine_odd_even Podd Peven n →

oddb n = false →

Peven n.

Proof.

(* FILL IN HERE *) Admitted.

☐

# Applying Theorems to Arguments

*proofs*as first-class objects.

Coq checks the
The reason is that the identifier plus_comm actually refers to a
Intuitively, this makes sense because the statement of a
theorem tells us what we can use that theorem for, just as the
type of a "computational" object tells us what we can do with that
object -- e.g., if we have a term of type nat → nat → nat, we
can give it two nats as arguments and get a nat back.
Similarly, if we have an object of type n = m → n + n = m + m
and we provide it an "argument" of type n = m, we can derive
n + n = m + m.
Operationally, this analogy goes even further: by applying a
theorem as if it were a function, i.e., applying it to hypotheses
with matching types, we can specialize its result without having
to resort to intermediate assertions. For example, suppose we
wanted to prove the following result:

*statement*of the plus_comm theorem (or prints it for us, if we leave off the part beginning with the colon) in the same way that it checks the*type*of any term that we ask it to Check. Why?*proof object*, which represents a logical derivation establishing of the truth of the statement ∀ n m : nat, n + m = m + n. The type of this object is the proposition which it is a proof of.
It appears at first sight that we ought to be able to prove this
by rewriting with plus_comm twice to make the two sides match.
The problem, however, is that the second rewrite will undo the
effect of the first.

Proof.

(* WORKED IN CLASS *)

intros x y z.

rewrite plus_comm.

rewrite plus_comm.

(* We are back where we started... *)

Abort.

We saw similar problems back in Chapter Induction, and saw one
way to work around them by using assert to derive a specialized
version of plus_comm that can be used to rewrite exactly where
we want.

Lemma plus_comm3_take2 :

∀ x y z, x + (y + z) = (z + y) + x.

Proof.

intros x y z.

rewrite plus_comm.

assert (H : y + z = z + y).

{ rewrite plus_comm. reflexivity. }

rewrite H.

reflexivity.

Qed.

A more elegant alternative is to apply plus_comm directly
to the arguments we want to instantiate it with, in much the same
way as we apply a polymorphic function to a type argument.

Lemma plus_comm3_take3 :

∀ x y z, x + (y + z) = (z + y) + x.

Proof.

intros x y z.

rewrite plus_comm.

rewrite (plus_comm y z).

reflexivity.

Qed.

Let's see another example of using a theorem like a
function. The following theorem says: any list l
containing some element must be nonempty.

Theorem in_not_nil :

∀ A (x : A) (l : list A), In x l → l ≠ [].

Proof.

intros A x l H. unfold not. intro Hl.

rewrite Hl in H.

simpl in H.

apply H.

Qed.

What makes this interesting is that one quantified variable
(x) does not appear in the conclusion (l ≠ []).
We should be able to use this theorem to prove the special case
where x is 42. However, naively, the tactic apply in_not_nil
will fail because it cannot infer the value of x.

Lemma in_not_nil_42 :

∀ l : list nat, In 42 l → l ≠ [].

Proof.

intros l H.

Fail apply in_not_nil.

Abort.

There are several ways to work around this:
Use apply ... with ...

Lemma in_not_nil_42_take2 :

∀ l : list nat, In 42 l → l ≠ [].

Proof.

intros l H.

apply in_not_nil with (x := 42).

apply H.

Qed.

∀ l : list nat, In 42 l → l ≠ [].

Proof.

intros l H.

apply in_not_nil with (x := 42).

apply H.

Qed.

Use apply ... in ...

Lemma in_not_nil_42_take3 :

∀ l : list nat, In 42 l → l ≠ [].

Proof.

intros l H.

apply in_not_nil in H.

apply H.

Qed.

∀ l : list nat, In 42 l → l ≠ [].

Proof.

intros l H.

apply in_not_nil in H.

apply H.

Qed.

Explicitly apply the lemma to the value for x.

Lemma in_not_nil_42_take4 :

∀ l : list nat, In 42 l → l ≠ [].

Proof.

intros l H.

apply (in_not_nil nat 42).

apply H.

Qed.

∀ l : list nat, In 42 l → l ≠ [].

Proof.

intros l H.

apply (in_not_nil nat 42).

apply H.

Qed.

Explicitly apply the lemma to a hypothesis.

Lemma in_not_nil_42_take5 :

∀ l : list nat, In 42 l → l ≠ [].

Proof.

intros l H.

apply (in_not_nil _ _ _ H).

Qed.

∀ l : list nat, In 42 l → l ≠ [].

Proof.

intros l H.

apply (in_not_nil _ _ _ H).

Qed.

You can "use theorems as functions" in this way with almost all
tactics that take a theorem name as an argument. Note also that
theorem application uses the same inference mechanisms as function
application; thus, it is possible, for example, to supply
wildcards as arguments to be inferred, or to declare some
hypotheses to a theorem as implicit by default. These features
are illustrated in the proof below. (The details of how this proof
works are not critical -- the goal here is just to illustrate what
can be done.)

Example lemma_application_ex :

∀ {n : nat} {ns : list nat},

In n (map (fun m ⇒ m × 0) ns) →

n = 0.

Proof.

intros n ns H.

destruct (proj1 _ _ (In_map_iff _ _ _ _ _) H)

as [m [Hm _]].

rewrite mult_0_r in Hm. rewrite <- Hm. reflexivity.

Qed.

We will see many more examples in later chapters.

# Coq vs. Set Theory

*Calculus of Inductive Constructions*, differs in some important ways from other formal systems that are used by mathematicians to write down precise and rigorous definitions and proofs. For example, in the most popular foundation for paper-and-pencil mathematics, Zermelo-Fraenkel Set Theory (ZFC), a mathematical object can potentially be a member of many different sets; a term in Coq's logic, on the other hand, is a member of at most one type. This difference often leads to slightly different ways of capturing informal mathematical concepts, but these are, by and large, about equally natural and easy to work with. For example, instead of saying that a natural number n belongs to the set of even numbers, we would say in Coq that even n holds, where even : nat → Prop is a property describing even numbers.

## Functional Extensionality

*any*type -- in particular, we can write propositions claiming that two

*functions*are equal to each other:

Example function_equality_ex

_{1}:

(fun x ⇒ 3 + x) = (fun x ⇒ (pred 4) + x).

Proof. reflexivity. Qed.

In common mathematical practice, two functions f and g are
considered equal if they produce the same output on every input:

(∀ x, f x = g x) → f = g
This is known as the principle of
Informally, an "extensional property" is one that pertains to an
object's observable behavior. Thus, functional extensionality
simply means that a function's identity is completely determined
by what we can observe from it -- i.e., the results we obtain
after applying it.
However, functional extensionality is not part of Coq's built-in
logic. This means that some apparently "obvious" propositions are
not provable.

(∀ x, f x = g x) → f = g

*functional extensionality*.Example function_equality_ex

_{2}:

(fun x ⇒ plus x 1) = (fun x ⇒ plus 1 x).

Proof.

(* Stuck *)

Abort.

However, we can add functional extensionality to Coq's core using
the Axiom command.

Defining something as an Axiom has the same effect as stating a
theorem and skipping its proof using Admitted, but it alerts the
reader that this isn't just something we're going to come back and
fill in later!
We can now invoke functional extensionality in proofs:

Example function_equality_ex

_{2}:

(fun x ⇒ plus x 1) = (fun x ⇒ plus 1 x).

Proof.

apply functional_extensionality. intros x.

apply plus_comm.

Qed.

Naturally, we must be careful when adding new axioms into Coq's
logic, as this can render it
Unfortunately, there is no simple way of telling whether an axiom
is safe to add: hard work by highly trained mathematicians is
often required to establish the consistency of any particular
combination of axioms.
Fortunately, it is known that adding functional extensionality, in
particular,
To check whether a particular proof relies on any additional
axioms, use the Print Assumptions command.

*inconsistent*-- that is, it may become possible to prove every proposition, including False, 2+2=5, etc.!*is*consistent.Print Assumptions function_equality_ex

_{2}.

(* ===>

Axioms:

functional_extensionality :

forall (X Y : Type) (f g : X -> Y),

(forall x : X, f x = g x) -> f = g *)

#### Exercise: 4 stars, standard (tr_rev_correct)

One problem with the definition of the list-reversing function rev that we have is that it performs a call to app on each step; running app takes time asymptotically linear in the size of the list, which means that rev is asymptotically quadratic. We can improve this with the following definitions:Fixpoint rev_append {X} (l

_{1}l

_{2}: list X) : list X :=

match l

_{1}with

| [] ⇒ l

_{2}

| x :: l

_{1}' ⇒ rev_append l

_{1}' (x :: l

_{2})

end.

Definition tr_rev {X} (l : list X) : list X :=

rev_append l [].

This version of rev is said to be
Prove that the two definitions are indeed equivalent.

*tail-recursive*, because the recursive call to the function is the last operation that needs to be performed (i.e., we don't have to execute ++ after the recursive call); a decent compiler will generate very efficient code in this case.## Propositions vs. Booleans

*booleans*(of type bool), and with

*propositions*(of type Prop).

... or that there exists some k such that n = double k.

Of course, it would be pretty strange if these two
characterizations of evenness did not describe the same set of
natural numbers! Fortunately, we can prove that they do...
We first need two helper lemmas.

Lemma evenb_double : ∀ k, evenb (double k) = true.

Proof.

intros k. induction k as [|k' IHk'].

- reflexivity.

- simpl. apply IHk'.

Qed.

intros k. induction k as [|k' IHk'].

- reflexivity.

- simpl. apply IHk'.

Qed.

Lemma evenb_double_conv : ∀ n, ∃ k,

n = if evenb n then double k else S (double k).

n = if evenb n then double k else S (double k).

Proof.

(* Hint: Use the evenb_S lemma from Induction.v. *)

(* FILL IN HERE *) Admitted.

☐

(* Hint: Use the evenb_S lemma from Induction.v. *)

(* FILL IN HERE *) Admitted.

☐

Now the main theorem:

Theorem even_bool_prop : ∀ n,

evenb n = true ↔ even n.

evenb n = true ↔ even n.

Proof.

intros n. split.

- intros H. destruct (evenb_double_conv n) as [k Hk].

rewrite Hk. rewrite H. ∃ k. reflexivity.

- intros [k Hk]. rewrite Hk. apply evenb_double.

Qed.

intros n. split.

- intros H. destruct (evenb_double_conv n) as [k Hk].

rewrite Hk. rewrite H. ∃ k. reflexivity.

- intros [k Hk]. rewrite Hk. apply evenb_double.

Qed.

In view of this theorem, we say that the boolean computation
evenb n is
Similarly, to state that two numbers n and m are equal, we can
say either

*reflected*in the truth of the proposition ∃ k, n = double k.- (1) that n =? m returns true, or
- (2) that n = m.

However, even when the boolean and propositional formulations of a
claim are equivalent from a purely logical perspective, they are
often not equivalent from the point of view of convenience for
some specific purpose.
In the case of even numbers above, when proving the
backwards direction of even_bool_prop (i.e., evenb_double,
going from the propositional to the boolean claim), we used a
simple induction on k. On the other hand, the converse (the
evenb_double_conv exercise) required a clever generalization,
since we can't directly prove (evenb n = true) → even n.
For these examples, the propositional claims are more useful than
their boolean counterparts, but this is not always the case. For
instance, we cannot test whether a general proposition is true or
not in a function definition; as a consequence, the following code
fragment is rejected:

Coq complains that n = 2 has type Prop, while it expects
an element of bool (or some other inductive type with two
elements). The reason has to do with the
Beyond the fact that non-computable properties are impossible in
general to phrase as boolean computations, even many
Conversely, an important side benefit of stating facts using
booleans is enabling some proof automation through computation
with Coq terms, a technique known as
Consider the following statement:

*computational*nature of Coq's core language, which is designed so that every function it can express is computable and total. One reason for this is to allow the extraction of executable programs from Coq developments. As a consequence, Prop in Coq does*not*have a universal case analysis operation telling whether any given proposition is true or false, since such an operation would allow us to write non-computable functions.*computable*properties are easier to express using Prop than bool, since recursive function definitions in Coq are subject to significant restrictions. For instance, the next chapter shows how to define the property that a regular expression matches a given string using Prop. Doing the same with bool would amount to writing a regular expression matching algorithm, which would be more complicated, harder to understand, and harder to reason about than a simple (non-algorithmic) definition of this property.*proof by reflection*.
The most direct way to prove this is to give the value of k
explicitly.

The proof of the corresponding boolean statement is even
simpler (because we don't have to invent the witness: Coq's
computation mechanism does it for us!).

What is interesting is that, since the two notions are equivalent,
we can use the boolean formulation to prove the other one without
mentioning the value 500 explicitly:

Although we haven't gained much in terms of proof-script
size in this case, larger proofs can often be made considerably
simpler by the use of reflection. As an extreme example, a famous
Coq proof of the even more famous
Another notable difference is that the negation of a "boolean
fact" is straightforward to state and prove: simply flip the
expected boolean result.

*4-color theorem*uses reflection to reduce the analysis of hundreds of different cases to a boolean computation.
In contrast, propositional negation can be more difficult
to work with directly.

Example not_even_1001' : ~(even 1001).

Proof.

(* WORKED IN CLASS *)

rewrite <- even_bool_prop.

unfold not.

simpl.

intro H.

discriminate H.

Qed.

Equality provides a complementary example, where it is sometimes
easier to work in the propositional world. Knowing that
n =? m = true is generally of little direct help in the middle of a
proof involving n and m; however, if we convert the statement
to the equivalent form n = m, we can rewrite with it.

Lemma plus_eqb_example : ∀ n m p : nat,

n =? m = true → n + p =? m + p = true.

Proof.

(* WORKED IN CLASS *)

intros n m p H.

rewrite eqb_eq in H.

rewrite H.

rewrite eqb_eq.

reflexivity.

Qed.

We won't discuss reflection any further here, but it serves as a
good example showing the complementary strengths of booleans and
general propositions, and being able to cross back and forth
between the boolean and propositional worlds will often be
convenient in later chapters.

#### Exercise: 2 stars, standard (logical_connectives)

The following theorems relate the propositional connectives studied in this chapter to the corresponding boolean operations.Theorem andb_true_iff : ∀ b

_{1}b

_{2}:bool,

b

_{1}&& b

_{2}= true ↔ b

_{1}= true ∧ b

_{2}= true.

Proof.

(* FILL IN HERE *) Admitted.

Theorem orb_true_iff : ∀ b

_{1}b

_{2},

b

_{1}|| b

_{2}= true ↔ b

_{1}= true ∨ b

_{2}= true.

Proof.

(* FILL IN HERE *) Admitted.

☐

#### Exercise: 1 star, standard (eqb_neq)

The following theorem is an alternate "negative" formulation of eqb_eq that is more convenient in certain situations. (We'll see examples in later chapters.) Hint: not_true_iff_false.#### Exercise: 3 stars, standard (eqb_list)

Given a boolean operator eqb for testing equality of elements of some type A, we can define a function eqb_list for testing equality of lists with elements in A. Complete the definition of the eqb_list function below. To make sure that your definition is correct, prove the lemma eqb_list_true_iff.Fixpoint eqb_list {A : Type} (eqb : A → A → bool)

(l

_{1}l

_{2}: list A) : bool

(* REPLACE THIS LINE WITH ":= _your_definition_ ." *). Admitted.

Theorem eqb_list_true_iff :

∀ A (eqb : A → A → bool),

(∀ a

_{1}a

_{2}, eqb a

_{1}a

_{2}= true ↔ a

_{1}= a

_{2}) →

∀ l

_{1}l

_{2}, eqb_list eqb l

_{1}l

_{2}= true ↔ l

_{1}= l

_{2}.

Proof.

(* FILL IN HERE *) Admitted.

☐

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

Recall the function forallb, from the exercise forall_exists_challenge in chapter Tactics:Fixpoint forallb {X : Type} (test : X → bool) (l : list X) : bool :=

match l with

| [] ⇒ true

| x :: l' ⇒ andb (test x) (forallb test l')

end.

Prove the theorem below, which relates forallb to the All
property defined above.

Theorem forallb_true_iff : ∀ X test (l : list X),

forallb test l = true ↔ All (fun x ⇒ test x = true) l.

Proof.

(* FILL IN HERE *) Admitted.

(Ungraded thought question) Are there any important properties of
the function forallb which are not captured by this
specification?

(* FILL IN HERE *)

☐

## Classical vs. Constructive Logic

*proofs*! In other words, the following intuitive reasoning principle is not derivable in Coq:

To understand operationally why this is the case, recall
that, to prove a statement of the form P ∨ Q, we use the left
and right tactics, which effectively require knowing which side
of the disjunction holds. But the universally quantified P in
excluded_middle is an
However, if we happen to know that P is reflected in some
boolean term b, then knowing whether it holds or not is trivial:
we just have to check the value of b.

*arbitrary*proposition, which we know nothing about. We don't have enough information to choose which of left or right to apply, just as Coq doesn't have enough information to mechanically decide whether P holds or not inside a function.Theorem restricted_excluded_middle : ∀ P b,

(P ↔ b = true) → P ∨ ¬ P.

Proof.

intros P [] H.

- left. rewrite H. reflexivity.

- right. rewrite H. intros contra. discriminate contra.

Qed.

In particular, the excluded middle is valid for equations n = m,
between natural numbers n and m.

Theorem restricted_excluded_middle_eq : ∀ (n m : nat),

n = m ∨ n ≠ m.

Proof.

intros n m.

apply (restricted_excluded_middle (n = m) (n =? m)).

symmetry.

apply eqb_eq.

Qed.

It may seem strange that the general excluded middle is not
available by default in Coq, since it is a standard feature of
familiar logics like ZFC. But there is a distinct advantage in
not assuming the excluded middle: statements in Coq make stronger
claims than the analogous statements in standard mathematics.
Notably, when there is a Coq proof of ∃ x, P x, it is
always possible to explicitly exhibit a value of x for which we
can prove P x -- in other words, every proof of existence is
Logics like Coq's, which do not assume the excluded middle, are
referred to as
More conventional logical systems such as ZFC, in which the
excluded middle does hold for arbitrary propositions, are referred
to as
The following example illustrates why assuming the excluded middle
may lead to non-constructive proofs:
Do you see what happened here? We used the excluded middle to
consider separately the cases where sqrt 2 ^ sqrt 2 is rational
and where it is not, without knowing which one actually holds!
Because of that, we finish the proof knowing that such a and b
exist but we cannot determine what their actual values are (at least,
not from this line of argument).
As useful as constructive logic is, it does have its limitations:
There are many statements that can easily be proven in classical
logic but that have only much more complicated constructive proofs, and
there are some that are known to have no constructive proof at
all! Fortunately, like functional extensionality, the excluded
middle is known to be compatible with Coq's logic, allowing us to
add it safely as an axiom. However, we will not need to do so
here: the results that we cover can be developed entirely
within constructive logic at negligible extra cost.
It takes some practice to understand which proof techniques must
be avoided in constructive reasoning, but arguments by
contradiction, in particular, are infamous for leading to
non-constructive proofs. Here's a typical example: suppose that
we want to show that there exists x with some property P,
i.e., such that P x. We start by assuming that our conclusion
is false; that is, ¬ ∃ x, P x. From this premise, it is not
hard to derive ∀ x, ¬ P x. If we manage to show that this
intermediate fact results in a contradiction, we arrive at an
existence proof without ever exhibiting a value of x for which
P x holds!
The technical flaw here, from a constructive standpoint, is that
we claimed to prove ∃ x, P x using a proof of
¬ ¬ (∃ x, P x). Allowing ourselves to remove double
negations from arbitrary statements is equivalent to assuming the
excluded middle, as shown in one of the exercises below. Thus,
this line of reasoning cannot be encoded in Coq without assuming
additional axioms.
Succinctly: for any proposition P,
Coq is consistent ==> (Coq + P ∨ ¬P) is consistent.
(Hint: you will need to come up with a clever assertion as the
next step in the proof.)

*constructive*.*constructive logics*.*classical*.*Claim*: There exist irrational numbers a and b such that a ^ b (a to the power b) is rational.*Proof*: It is not difficult to show that sqrt 2 is irrational. If sqrt 2 ^ sqrt 2 is rational, it suffices to take a = b = sqrt 2 and we are done. Otherwise, sqrt 2 ^ sqrt 2 is irrational. In this case, we can take a = sqrt 2 ^ sqrt 2 and b = sqrt 2, since a ^ b = sqrt 2 ^ (sqrt 2 × sqrt 2) = sqrt 2 ^ 2 = 2. ☐#### Exercise: 3 stars, standard (excluded_middle_irrefutable)

Proving the consistency of Coq with the general excluded middle axiom requires complicated reasoning that cannot be carried out within Coq itself. However, the following theorem implies that it is always safe to assume a decidability axiom (i.e., an instance of excluded middle) for any*particular*Prop P. Why? Because we cannot prove the negation of such an axiom. If we could, we would have both ¬ (P ∨ ¬P) and ¬ ¬ (P ∨ ¬P) (since P implies ¬ ¬ P, by lemma double_neg, which we proved above), which would be a contradiction. But since we can't, it is safe to add P ∨ ¬P as an axiom.Theorem excluded_middle_irrefutable: ∀ (P:Prop),

¬ ¬ (P ∨ ¬ P).

Proof.

unfold not. intros P H.

(* FILL IN HERE *) Admitted.

☐

#### Exercise: 3 stars, advanced (not_exists_dist)

It is a theorem of classical logic that the following two assertions are equivalent:¬(∃ x, ¬P x)

∀ x, P x

Theorem not_exists_dist :

excluded_middle →

∀ (X:Type) (P : X → Prop),

¬ (∃ x, ¬ P x) → (∀ x, P x).

Proof.

(* FILL IN HERE *) Admitted.

☐

#### Exercise: 5 stars, standard, optional (classical_axioms)

For those who like a challenge, here is an exercise taken from the Coq'Art book by Bertot and Casteran (p. 123). Each of the following four statements, together with excluded_middle, can be considered as characterizing classical logic. We can't prove any of them in Coq, but we can consistently add any one of them as an axiom if we wish to work in classical logic.Definition peirce := ∀ P Q: Prop,

((P→Q)->P)->P.

Definition double_negation_elimination := ∀ P:Prop,

~~P → P.

Definition de_morgan_not_and_not := ∀ P Q:Prop,

~(~P ∧ ¬Q) → P∨Q.

Definition implies_to_or := ∀ P Q:Prop,

(P→Q) → (¬P∨Q).

(* FILL IN HERE *)

☐

(* 2020-08-25 12:53 *)