MoreIndMore on Induction
Require Export ProofObjects.
Induction Principles
Check nat_ind.
(* ===> nat_ind :
forall P : nat -> Prop,
P 0 ->
(forall n : nat, P n -> P (S n)) ->
forall n : nat, P n *)
The induction tactic is a straightforward wrapper that, at
its core, simply performs apply t_ind. To see this more
clearly, let's experiment a little with using apply nat_ind
directly, instead of the induction tactic, to carry out some
proofs. Here, for example, is an alternate proof of a theorem
that we saw in the Basics chapter.
Theorem mult_0_r' : ∀n:nat,
n × 0 = 0.
Proof.
apply nat_ind.
Case "O". reflexivity.
Case "S". simpl. intros n IHn. rewrite → IHn.
reflexivity. Qed.
This proof is basically the same as the earlier one, but a
few minor differences are worth noting. First, in the induction
step of the proof (the "S" case), we have to do a little
bookkeeping manually (the intros) that induction does
automatically.
Second, we do not introduce n into the context before applying
nat_ind — the conclusion of nat_ind is a quantified formula,
and apply needs this conclusion to exactly match the shape of
the goal state, including the quantifier. The induction tactic
works either with a variable in the context or a quantified
variable in the goal.
Third, the apply tactic automatically chooses variable names for
us (in the second subgoal, here), whereas induction lets us
specify (with the as... clause) what names should be used. The
automatic choice is actually a little unfortunate, since it
re-uses the name n for a variable that is different from the n
in the original theorem. This is why the Case annotation is
just S — if we tried to write it out in the more explicit form
that we've been using for most proofs, we'd have to write n = S
n, which doesn't make a lot of sense! All of these conveniences
make induction nicer to use in practice than applying induction
principles like nat_ind directly. But it is important to
realize that, modulo this little bit of bookkeeping, applying
nat_ind is what we are really doing.
Exercise: 2 stars, optional (plus_one_r')
Complete this proof as we did mult_0_r' above, without using the induction tactic.Theorem plus_one_r' : ∀n:nat,
n + 1 = S n.
Proof.
(* FILL IN HERE *) Admitted.
☐
Coq generates induction principles for every datatype defined with
Inductive, including those that aren't recursive. (Although
we don't need induction to prove properties of non-recursive
datatypes, the idea of an induction principle still makes sense
for them: it gives a way to prove that a property holds for all
values of the type.)
These generated principles follow a similar pattern. If we define a
type t with constructors c1 ... cn, Coq generates a theorem
with this shape:
t_ind :
∀P : t → Prop,
... case for c1 ... →
... case for c2 ... →
...
... case for cn ... →
∀n : t, P n
The specific shape of each case depends on the arguments to the
corresponding constructor. Before trying to write down a general
rule, let's look at some more examples. First, an example where
the constructors take no arguments:
∀P : t → Prop,
... case for c1 ... →
... case for c2 ... →
...
... case for cn ... →
∀n : t, P n
Inductive yesno : Type :=
| yes : yesno
| no : yesno.
Check yesno_ind.
(* ===> yesno_ind : forall P : yesno -> Prop,
P yes ->
P no ->
forall y : yesno, P y *)
Exercise: 1 star, optional (rgb)
Write out the induction principle that Coq will generate for the following datatype. Write down your answer on paper or type it into a comment, and then compare it with what Coq prints.Inductive rgb : Type :=
| red : rgb
| green : rgb
| blue : rgb.
Check rgb_ind.
☐
Here's another example, this time with one of the constructors
taking some arguments.
Inductive natlist : Type :=
| nnil : natlist
| ncons : nat → natlist → natlist.
Check natlist_ind.
(* ===> (modulo a little variable renaming for clarity)
natlist_ind :
forall P : natlist -> Prop,
P nnil ->
(forall (n : nat) (l : natlist), P l -> P (ncons n l)) ->
forall n : natlist, P n *)
Exercise: 1 star, optional (natlist1)
Suppose we had written the above definition a little differently:Inductive natlist1 : Type :=
| nnil1 : natlist1
| nsnoc1 : natlist1 → nat → natlist1.
Now what will the induction principle look like? ☐
From these examples, we can extract this general rule:
- The type declaration gives several constructors; each corresponds to one clause of the induction principle.
- Each constructor c takes argument types a1...an.
- Each ai can be either t (the datatype we are defining) or some other type s.
- The corresponding case of the induction principle
says (in English):
- "for all values x1...xn of types a1...an, if P holds for each of the inductive arguments (each xi of type t), then P holds for c x1 ... xn".
Exercise: 1 star, optional (byntree_ind)
Write out the induction principle that Coq will generate for the following datatype. Write down your answer on paper or type it into a comment, and then compare it with what Coq prints.Inductive byntree : Type :=
| bempty : byntree
| bleaf : yesno → byntree
| nbranch : yesno → byntree → byntree → byntree.
☐
Exercise: 1 star, optional (ex_set)
Here is an induction principle for an inductively defined set.
ExSet_ind :
∀P : ExSet → Prop,
(∀b : bool, P (con1 b)) →
(∀(n : nat) (e : ExSet), P e → P (con2 n e)) →
∀e : ExSet, P e
Give an Inductive definition of ExSet:
∀P : ExSet → Prop,
(∀b : bool, P (con1 b)) →
(∀(n : nat) (e : ExSet), P e → P (con2 n e)) →
∀e : ExSet, P e
Inductive ExSet : Type :=
(* FILL IN HERE *)
.
☐
What about polymorphic datatypes?
The inductive definition of polymorphic lists
Inductive list (X:Type) : Type :=
| nil : list X
| cons : X → list X → list X.
is very similar to that of natlist. The main difference is
that, here, the whole definition is parameterized on a set X:
that is, we are defining a family of inductive types list X,
one for each X. (Note that, wherever list appears in the body
of the declaration, it is always applied to the parameter X.)
The induction principle is likewise parameterized on X:
| nil : list X
| cons : X → list X → list X.
list_ind :
∀(X : Type) (P : list X → Prop),
P [] →
(∀(x : X) (l : list X), P l → P (x :: l)) →
∀l : list X, P l
Note the wording here (and, accordingly, the form of list_ind):
The whole induction principle is parameterized on X. That is,
list_ind can be thought of as a polymorphic function that, when
applied to a type X, gives us back an induction principle
specialized to the type list X.
∀(X : Type) (P : list X → Prop),
P [] →
(∀(x : X) (l : list X), P l → P (x :: l)) →
∀l : list X, P l
Exercise: 1 star, optional (tree)
Write out the induction principle that Coq will generate for the following datatype. Compare your answer with what Coq prints.Inductive tree (X:Type) : Type :=
| leaf : X → tree X
| node : tree X → tree X → tree X.
Check tree_ind.
☐
Exercise: 1 star, optional (mytype)
Find an inductive definition that gives rise to the following induction principle:
mytype_ind :
∀(X : Type) (P : mytype X → Prop),
(∀x : X, P (constr1 X x)) →
(∀n : nat, P (constr2 X n)) →
(∀m : mytype X, P m →
∀n : nat, P (constr3 X m n)) →
∀m : mytype X, P m
☐
∀(X : Type) (P : mytype X → Prop),
(∀x : X, P (constr1 X x)) →
(∀n : nat, P (constr2 X n)) →
(∀m : mytype X, P m →
∀n : nat, P (constr3 X m n)) →
∀m : mytype X, P m
Exercise: 1 star, optional (foo)
Find an inductive definition that gives rise to the following induction principle:
foo_ind :
∀(X Y : Type) (P : foo X Y → Prop),
(∀x : X, P (bar X Y x)) →
(∀y : Y, P (baz X Y y)) →
(∀f1 : nat → foo X Y,
(∀n : nat, P (f1 n)) → P (quux X Y f1)) →
∀f2 : foo X Y, P f2
☐
∀(X Y : Type) (P : foo X Y → Prop),
(∀x : X, P (bar X Y x)) →
(∀y : Y, P (baz X Y y)) →
(∀f1 : nat → foo X Y,
(∀n : nat, P (f1 n)) → P (quux X Y f1)) →
∀f2 : foo X Y, P f2
Exercise: 1 star, optional (foo')
Consider the following inductive definition:Inductive foo' (X:Type) : Type :=
| C1 : list X → foo' X → foo' X
| C2 : foo' X.
What induction principle will Coq generate for foo'? Fill
in the blanks, then check your answer with Coq.)
☐
foo'_ind :
∀(X : Type) (P : foo' X → Prop),
(∀(l : list X) (f : foo' X),
_______________________ →
_______________________ ) →
___________________________________________ →
∀f : foo' X, ________________________
∀(X : Type) (P : foo' X → Prop),
(∀(l : list X) (f : foo' X),
_______________________ →
_______________________ ) →
___________________________________________ →
∀f : foo' X, ________________________
Induction Hypotheses
∀P : nat → Prop,
P 0 →
(∀n : nat, P n → P (S n)) →
∀n : nat, P n
is a generic statement that holds for all propositions
P (strictly speaking, for all families of propositions P
indexed by a number n). Each time we use this principle, we
are choosing P to be a particular expression of type
nat→Prop.
P 0 →
(∀n : nat, P n → P (S n)) →
∀n : nat, P n
Definition P_m0r (n:nat) : Prop :=
n × 0 = 0.
... or equivalently...
Definition P_m0r' : nat→Prop :=
fun n ⇒ n × 0 = 0.
Now when we do the proof it is easier to see where P_m0r
appears.
Theorem mult_0_r'' : ∀n:nat,
P_m0r n.
Proof.
apply nat_ind.
Case "n = O". reflexivity.
Case "n = S n'".
(* Note the proof state at this point! *)
intros n IHn.
unfold P_m0r in IHn. unfold P_m0r. simpl. apply IHn. Qed.
This extra naming step isn't something that we'll do in
normal proofs, but it is useful to do it explicitly for an example
or two, because it allows us to see exactly what the induction
hypothesis is. If we prove ∀ n, P_m0r n by induction on
n (using either induction or apply nat_ind), we see that the
first subgoal requires us to prove P_m0r 0 ("P holds for
zero"), while the second subgoal requires us to prove ∀ n',
P_m0r n' → P_m0r n' (S n') (that is "P holds of S n' if it
holds of n'" or, more elegantly, "P is preserved by S").
The induction hypothesis is the premise of this latter
implication — the assumption that P holds of n', which we are
allowed to use in proving that P holds for S n'.
More on the induction Tactic
- If P n is some proposition involving a natural number n, and
we want to show that P holds for all numbers n, we can
reason like this:
- show that P O holds
- show that, if P n' holds, then so does P (S n')
- conclude that P n holds for all n.
Theorem plus_assoc' : ∀n m p : nat,
n + (m + p) = (n + m) + p.
Proof.
(* ...we first introduce all 3 variables into the context,
which amounts to saying "Consider an arbitrary n, m, and
p..." *)
intros n m p.
(* ...We now use the induction tactic to prove P n (that
is, n + (m + p) = (n + m) + p) for _all_ n,
and hence also for the particular n that is in the context
at the moment. *)
induction n as [| n'].
Case "n = O". reflexivity.
Case "n = S n'".
(* In the second subgoal generated by induction -- the
"inductive step" -- we must prove that P n' implies
P (S n') for all n'. The induction tactic
automatically introduces n' and P n' into the context
for us, leaving just P (S n') as the goal. *)
simpl. rewrite → IHn'. reflexivity. Qed.
It also works to apply induction to a variable that is
quantified in the goal.
Theorem plus_comm' : ∀n m : nat,
n + m = m + n.
Proof.
induction n as [| n'].
Case "n = O". intros m. rewrite → plus_0_r. reflexivity.
Case "n = S n'". intros m. simpl. rewrite → IHn'.
rewrite ← plus_n_Sm. reflexivity. Qed.
Note that induction n leaves m still bound in the goal —
i.e., what we are proving inductively is a statement beginning
with ∀ m.
If we do induction on a variable that is quantified in the goal
after some other quantifiers, the induction tactic will
automatically introduce the variables bound by these quantifiers
into the context.
Theorem plus_comm'' : ∀n m : nat,
n + m = m + n.
Proof.
(* Let's do induction on m this time, instead of n... *)
induction m as [| m'].
Case "m = O". simpl. rewrite → plus_0_r. reflexivity.
Case "m = S m'". simpl. rewrite ← IHm'.
rewrite ← plus_n_Sm. reflexivity. Qed.
Exercise: 1 star, optional (plus_explicit_prop)
Rewrite both plus_assoc' and plus_comm' and their proofs in the same style as mult_0_r'' above — that is, for each theorem, give an explicit Definition of the proposition being proved by induction, and state the theorem and proof in terms of this defined proposition.(* FILL IN HERE *)
☐
One potentially confusing feature of the induction tactic is
that it happily lets you try to set up an induction over a term
that isn't sufficiently general. The net effect of this will be
do lose information (much as destruct can do), and leave
you unable to complete the proof. Here's an example:
Generalizing Inductions.
Lemma one_not_beautiful_FAILED: ¬ beautiful 1.
Proof.
intro H.
(* Just doing an inversion on H won't get us very far in the b_sum
case. (Try it!). So we'll need induction. A naive first attempt: *)
induction H.
(* But now, although we get four cases, as we would expect from
the definition of beautiful, we lose all information about H ! *)
Abort.
The problem is that induction over a Prop only works properly over
completely general instances of the Prop, i.e. one in which all
the arguments are free (unconstrained) variables.
In this respect it behaves more
like destruct than like inversion.
When you're tempted to do use induction like this, it is generally
an indication that you need to be proving something more general.
But in some cases, it suffices to pull out any concrete arguments
into separate equations, like this:
Lemma one_not_beautiful: ∀n, n = 1 → ¬ beautiful n.
Proof.
intros n E H.
induction H as [| | | p q Hp IHp Hq IHq].
Case "b_0".
inversion E.
Case "b_3".
inversion E.
Case "b_5".
inversion E.
Case "b_sum".
(* the rest is a tedious case analysis *)
destruct p as [|p'].
SCase "p = 0".
destruct q as [|q'].
SSCase "q = 0".
inversion E.
SSCase "q = S q'".
apply IHq. apply E.
SCase "p = S p'".
destruct q as [|q'].
SSCase "q = 0".
apply IHp. rewrite plus_0_r in E. apply E.
SSCase "q = S q'".
simpl in E. inversion E. destruct p'. inversion H0. inversion H0.
Qed.
There's a handy remember tactic that can generate the second
proof state out of the original one.
Lemma one_not_beautiful': ¬ beautiful 1.
Proof.
intros H.
remember 1 as n eqn:E.
(* now carry on as above *)
induction H.
Admitted.
Informal Proofs (Advanced)
Informal Proofs by Induction
Induction Over an Inductively Defined Set
- Theorem: <Universally quantified proposition of the form
"For all n:S, P(n)," where S is some inductively defined
set.>
- Suppose n = c a1 ... ak, where <...and here we state
the IH for each of the a's that has type S, if any>.
We must show <...and here we restate P(c a1 ... ak)>.
- <other cases similarly...> ☐
- Suppose n = c a1 ... ak, where <...and here we state
the IH for each of the a's that has type S, if any>.
We must show <...and here we restate P(c a1 ... ak)>.
- Theorem: For all sets X, lists l : list X, and numbers
n, if length l = n then index (S n) l = None.
- Suppose l = []. We must show, for all numbers n,
that, if length [] = n, then index (S n) [] =
None.
- Suppose l = x :: l' for some x and l', where
length l' = n' implies index (S n') l' = None, for
any number n'. We must show, for all n, that, if
length (x::l') = n then index (S n) (x::l') =
None.
length l = length (x::l') = S (length l'),it suffices to show thatindex (S (length l')) l' = None.But this follows directly from the induction hypothesis, picking n' to be length l'. ☐
- Suppose l = []. We must show, for all numbers n,
that, if length [] = n, then index (S n) [] =
None.
Induction Over an Inductively Defined Proposition
- Theorem: <Proposition of the form "Q → P," where Q is
some inductively defined proposition (more generally,
"For all x y z, Q x y z → P x y z")>
- Suppose the final rule used to show Q is c. Then
<...and here we state the types of all of the a's
together with any equalities that follow from the
definition of the constructor and the IH for each of
the a's that has type Q, if there are any>. We must
show <...and here we restate P>.
- <other cases similarly...> ☐
- Suppose the final rule used to show Q is c. Then
<...and here we state the types of all of the a's
together with any equalities that follow from the
definition of the constructor and the IH for each of
the a's that has type Q, if there are any>. We must
show <...and here we restate P>.
- Theorem: The ≤ relation is transitive — i.e., for all
numbers n, m, and o, if n ≤ m and m ≤ o, then
n ≤ o.
- Suppose the final rule used to show m ≤ o is
le_n. Then m = o and we must show that n ≤ m,
which is immediate by hypothesis.
- Suppose the final rule used to show m ≤ o is
le_S. Then o = S o' for some o' with m ≤ o'.
We must show that n ≤ S o'.
By induction hypothesis, n ≤ o'.
- Suppose the final rule used to show m ≤ o is
le_n. Then m = o and we must show that n ≤ m,
which is immediate by hypothesis.
Optional Material
Induction Principles in Prop
Inductive gorgeous : nat → Prop :=
g_0 : gorgeous 0
| g_plus3 : ∀n, gorgeous n → gorgeous (3+m)
| g_plus5 : ∀n, gorgeous n → gorgeous (5+m).
...to give rise to an induction principle that looks like this...
g_0 : gorgeous 0
| g_plus3 : ∀n, gorgeous n → gorgeous (3+m)
| g_plus5 : ∀n, gorgeous n → gorgeous (5+m).
gorgeous_ind_max :
∀P : (∀n : nat, gorgeous n → Prop),
P O g_0 →
(∀(m : nat) (e : gorgeous m),
P m e → P (3+m) (g_plus3 m e) →
(∀(m : nat) (e : gorgeous m),
P m e → P (5+m) (g_plus5 m e) →
∀(n : nat) (e : gorgeous n), P n e
... because:
∀P : (∀n : nat, gorgeous n → Prop),
P O g_0 →
(∀(m : nat) (e : gorgeous m),
P m e → P (3+m) (g_plus3 m e) →
(∀(m : nat) (e : gorgeous m),
P m e → P (5+m) (g_plus5 m e) →
∀(n : nat) (e : gorgeous n), P n e
- Since gorgeous is indexed by a number n (every gorgeous
object e is a piece of evidence that some particular number
n is gorgeous), the proposition P is parameterized by both
n and e — that is, the induction principle can be used to
prove assertions involving both a gorgeous number and the
evidence that it is gorgeous.
- Since there are three ways of giving evidence of gorgeousness
(gorgeous has three constructors), applying the induction
principle generates three subgoals:
- We must prove that P holds for O and b_0.
- We must prove that, whenever n is a gorgeous
number and e is an evidence of its gorgeousness,
if P holds of n and e,
then it also holds of 3+m and g_plus3 n e.
- We must prove that, whenever n is a gorgeous
number and e is an evidence of its gorgeousness,
if P holds of n and e,
then it also holds of 5+m and g_plus5 n e.
- We must prove that P holds for O and b_0.
- If these subgoals can be proved, then the induction principle tells us that P is true for all gorgeous numbers n and evidence e of their gorgeousness.
∀P : nat → Prop,
... →
∀n : nat, gorgeous n → P n
For this reason, Coq actually generates the following simplified
induction principle for gorgeous:
... →
∀n : nat, gorgeous n → P n
Check gorgeous_ind.
(* ===> gorgeous_ind
: forall P : nat -> Prop,
P 0 ->
(forall n : nat, gorgeous n -> P n -> P (3 + n)) ->
(forall n : nat, gorgeous n -> P n -> P (5 + n)) ->
forall n : nat, gorgeous n -> P n *)
In particular, Coq has dropped the evidence term e as a
parameter of the the proposition P, and consequently has
rewritten the assumption ∀ (n : nat) (e: gorgeous n), ...
to be ∀ (n : nat), gorgeous n → ...; i.e., we no longer
require explicit evidence of the provability of gorgeous n.
In English, gorgeous_ind says:
As expected, we can apply gorgeous_ind directly instead of using induction.
- Suppose, P is a property of natural numbers (that is, P n is
a Prop for every n). To show that P n holds whenever n
is gorgeous, it suffices to show:
- P holds for 0,
- for any n, if n is gorgeous and P holds for
n, then P holds for 3+n,
- for any n, if n is gorgeous and P holds for n, then P holds for 5+n.
- P holds for 0,
Theorem gorgeous__beautiful' : ∀n, gorgeous n → beautiful n.
Proof.
intros.
apply gorgeous_ind.
Case "g_0".
apply b_0.
Case "g_plus3".
intros.
apply b_sum. apply b_3.
apply H1.
Case "g_plus5".
intros.
apply b_sum. apply b_5.
apply H1.
apply H.
Qed.
The precise form of an Inductive definition can affect the
induction principle Coq generates.
For example, in Logic, we have defined ≤ as:
(* Inductive le : nat -> nat -> Prop :=
| le_n : forall n, le n n
| le_S : forall n m, (le n m) -> (le n (S m)). *)
This definition can be streamlined a little by observing that the
left-hand argument n is the same everywhere in the definition,
so we can actually make it a "general parameter" to the whole
definition, rather than an argument to each constructor.
Inductive le (n:nat) : nat → Prop :=
| le_n : le n n
| le_S : ∀m, (le n m) → (le n (S m)).
Notation "m ≤ n" := (le m n).
The second one is better, even though it looks less symmetric.
Why? Because it gives us a simpler induction principle.
Check 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 *)
By contrast, the induction principle that Coq calculates for the
first definition has a lot of extra quantifiers, which makes it
messier to work with when proving things by induction. Here is
the induction principle for the first le:
(* le_ind :
forall P : nat -> nat -> Prop,
(forall n : nat, P n n) ->
(forall n m : nat, le n m -> P n m -> P n (S m)) ->
forall n n0 : nat, le n n0 -> P n n0 *)
Additional Exercises
Exercise: 2 stars, optional (foo_ind_principle)
Suppose we make the following inductive definition:
Inductive foo (X : Set) (Y : Set) : Set :=
| foo1 : X → foo X Y
| foo2 : Y → foo X Y
| foo3 : foo X Y → foo X Y.
Fill in the blanks to complete the induction principle that will be
generated by Coq.
| foo1 : X → foo X Y
| foo2 : Y → foo X Y
| foo3 : foo X Y → foo X Y.
foo_ind
: ∀(X Y : Set) (P : foo X Y → Prop),
(∀x : X, __________________________________) →
(∀y : Y, __________________________________) →
(________________________________________________) →
________________________________________________
: ∀(X Y : Set) (P : foo X Y → Prop),
(∀x : X, __________________________________) →
(∀y : Y, __________________________________) →
(________________________________________________) →
________________________________________________
Exercise: 2 stars, optional (bar_ind_principle)
Consider the following induction principle:
bar_ind
: ∀P : bar → Prop,
(∀n : nat, P (bar1 n)) →
(∀b : bar, P b → P (bar2 b)) →
(∀(b : bool) (b0 : bar), P b0 → P (bar3 b b0)) →
∀b : bar, P b
Write out the corresponding inductive set definition.
: ∀P : bar → Prop,
(∀n : nat, P (bar1 n)) →
(∀b : bar, P b → P (bar2 b)) →
(∀(b : bool) (b0 : bar), P b0 → P (bar3 b b0)) →
∀b : bar, P b
Inductive bar : Set :=
| bar1 : ________________________________________
| bar2 : ________________________________________
| bar3 : ________________________________________.
| bar1 : ________________________________________
| bar2 : ________________________________________
| bar3 : ________________________________________.
Exercise: 2 stars, optional (no_longer_than_ind)
Given the following inductively defined proposition:
Inductive no_longer_than (X : Set) : (list X) → nat → Prop :=
| nlt_nil : ∀n, no_longer_than X [] n
| nlt_cons : ∀x l n, no_longer_than X l n →
no_longer_than X (x::l) (S n)
| nlt_succ : ∀l n, no_longer_than X l n →
no_longer_than X l (S n).
write the induction principle generated by Coq.
| nlt_nil : ∀n, no_longer_than X [] n
| nlt_cons : ∀x l n, no_longer_than X l n →
no_longer_than X (x::l) (S n)
| nlt_succ : ∀l n, no_longer_than X l n →
no_longer_than X l (S n).
no_longer_than_ind
: ∀(X : Set) (P : list X → nat → Prop),
(∀n : nat, ____________________) →
(∀(x : X) (l : list X) (n : nat),
no_longer_than X l n → ____________________ →
_____________________________ →
(∀(l : list X) (n : nat),
no_longer_than X l n → ____________________ →
_____________________________ →
∀(l : list X) (n : nat), no_longer_than X l n →
____________________
: ∀(X : Set) (P : list X → nat → Prop),
(∀n : nat, ____________________) →
(∀(x : X) (l : list X) (n : nat),
no_longer_than X l n → ____________________ →
_____________________________ →
(∀(l : list X) (n : nat),
no_longer_than X l n → ____________________ →
_____________________________ →
∀(l : list X) (n : nat), no_longer_than X l n →
____________________
(* Inductive eq (X:Type) : X -> X -> Prop :=
refl_equal : forall x, eq X x x. *)
In the Coq standard library, the definition of equality is
slightly different:
Inductive eq' (X:Type) (x:X) : X → Prop :=
refl_equal' : eq' X x x.
The advantage of this definition is that the induction
principle that Coq derives for it is precisely the familiar
principle of Leibniz equality: what we mean when we say "x and
y are equal" is that every property on P that is true of x
is also true of y.
Check eq'_ind.
(* ===>
forall (X : Type) (x : X) (P : X -> Prop),
P x -> forall y : X, x =' y -> P y
===> (i.e., after a little reorganization)
forall (X : Type) (x : X) forall y : X,
x =' y ->
forall P : X -> Prop, P x -> P y *)
The induction principles for conjunction and disjunction are a
good illustration of Coq's way of generating simplified induction
principles for Inductively defined propositions, which we
discussed above. You try first:
Exercise: 1 star, optional (and_ind_principle)
See if you can predict the induction principle for conjunction.(* Check and_ind. *)
☐
Exercise: 1 star, optional (or_ind_principle)
See if you can predict the induction principle for disjunction.(* Check or_ind. *)
☐
Check and_ind.
From the inductive definition of the proposition and P Q
Inductive and (P Q : Prop) : Prop :=
conj : P → Q → (and P Q).
we might expect Coq to generate this induction principle
conj : P → Q → (and P Q).
and_ind_max :
∀(P Q : Prop) (P0 : P ∧ Q → Prop),
(∀(a : P) (b : Q), P0 (conj P Q a b)) →
∀a : P ∧ Q, P0 a
but actually it generates this simpler and more useful one:
∀(P Q : Prop) (P0 : P ∧ Q → Prop),
(∀(a : P) (b : Q), P0 (conj P Q a b)) →
∀a : P ∧ Q, P0 a
and_ind :
∀P Q P0 : Prop,
(P → Q → P0) →
P ∧ Q → P0
In the same way, when given the inductive definition of or P Q
∀P Q P0 : Prop,
(P → Q → P0) →
P ∧ Q → P0
Inductive or (P Q : Prop) : Prop :=
| or_introl : P → or P Q
| or_intror : Q → or P Q.
instead of the "maximal induction principle"
| or_introl : P → or P Q
| or_intror : Q → or P Q.
or_ind_max :
∀(P Q : Prop) (P0 : P ∨ Q → Prop),
(∀a : P, P0 (or_introl P Q a)) →
(∀b : Q, P0 (or_intror P Q b)) →
∀o : P ∨ Q, P0 o
what Coq actually generates is this:
∀(P Q : Prop) (P0 : P ∨ Q → Prop),
(∀a : P, P0 (or_introl P Q a)) →
(∀b : Q, P0 (or_intror P Q b)) →
∀o : P ∨ Q, P0 o
or_ind :
∀P Q P0 : Prop,
(P → P0) →
(Q → P0) →
P ∨ Q → P0
∀P Q P0 : Prop,
(P → P0) →
(Q → P0) →
P ∨ Q → P0
Exercise: 1 star, optional (False_ind_principle)
Can you predict the induction principle for falsehood?(* Check False_ind. *)
☐
Here's the induction principle that Coq generates for existentials:
Check ex_ind.
(* ===> forall (X:Type) (P: X->Prop) (Q: Prop),
(forall witness:X, P witness -> Q) ->
ex X P ->
Q *)
This induction principle can be understood as follows: If we have
a function f that can construct evidence for Q given any
witness of type X together with evidence that this witness has
property P, then from a proof of ex X P we can extract the
witness and evidence that must have been supplied to the
constructor, give these to f, and thus obtain a proof of Q.
Explicit Proof Objects for Induction
Check nat_ind.
(* ===>
nat_ind : forall P : nat -> Prop,
P 0 ->
(forall n : nat, P n -> P (S n)) ->
forall n : nat, P n *)
There's nothing magic about this induction lemma: it's just
another Coq lemma that requires a proof. Coq generates the proof
automatically too...
Print nat_ind.
Print nat_rect.
(* ===> (after some manual inlining and tidying)
nat_ind =
fun (P : nat -> Prop)
(f : P 0)
(f0 : forall n : nat, P n -> P (S n)) =>
fix F (n : nat) : P n :=
match n with
| 0 => f
| S n0 => f0 n0 (F n0)
end.
*)
We can read this as follows:
Suppose we have evidence f that P holds on 0, and
evidence f0 that ∀ n:nat, P n → P (S n).
Then we can prove that P holds of an arbitrary nat n via
a recursive function F (here defined using the expression
form Fix rather than by a top-level Fixpoint
declaration). F pattern matches on n:
We can adapt this approach to proving nat_ind to help prove
non-standard induction principles too. Recall our desire to
prove that
∀ n : nat, even n → ev n.
Attempts to do this by standard induction on n fail, because the
induction principle only lets us proceed when we can prove that
even n → even (S n) — which is of course never provable. What
we did in Logic was a bit of a hack:
Theorem even__ev : ∀ n : nat,
(even n → ev n) ∧ (even (S n) → ev (S n)).
We can make a much better proof by defining and proving a
non-standard induction principle that goes "by twos":
- If it finds 0, F uses f to show that P n holds.
- If it finds S n0, F applies itself recursively on n0 to obtain evidence that P n0 holds; then it applies f0 on that evidence to show that P (S n) holds.
Definition nat_ind2 :
∀(P : nat → Prop),
P 0 →
P 1 →
(∀n : nat, P n → P (S(S n))) →
∀n : nat , P n :=
fun P ⇒ fun P0 ⇒ fun P1 ⇒ fun PSS ⇒
fix f (n:nat) := match n with
0 ⇒ P0
| 1 ⇒ P1
| S (S n') ⇒ PSS n' (f n')
end.
Once you get the hang of it, it is entirely straightforward to
give an explicit proof term for induction principles like this.
Proving this as a lemma using tactics is much less intuitive (try
it!).
The induction ... using tactic variant gives a convenient way to
specify a non-standard induction principle like this.
Lemma even__ev' : ∀n, even n → ev n.
Proof.
intros.
induction n as [ | |n'] using nat_ind2.
Case "even 0".
apply ev_0.
Case "even 1".
inversion H.
Case "even (S(S n'))".
apply ev_SS.
apply IHn'. unfold even. unfold even in H. simpl in H. apply H.
Qed.
The Coq Trusted Computing Base
- Since Coq types can themselves be expressions, the checker must
normalize these (by using the computation rules) before
comparing them.
- The checker must make sure that match expressions are
exhaustive. That is, there must be an arm for every possible
constructor. To see why, consider the following alleged proof
object:
Definition or_bogus : ∀P Q, P ∨ Q → P :=All the types here match correctly, but the match only considers one of the possible constructors for or. Coq's exhaustiveness check will reject this definition.
fun (P Q : Prop) (A : P ∨ Q) ⇒
match A with
| or_introl H ⇒ H
end. - The checker must make sure that each fix expression
terminates. It does this using a syntactic check to make sure
that each recursive call is on a subexpression of the original
argument. To see why this is essential, consider this alleged
proof:
Definition nat_false : ∀(n:nat), False :=Again, this is perfectly well-typed, but (fortunately) Coq will reject it.
fix f (n:nat) : False := f n.
(* $Date: 2013-07-17 16:19:11 -0400 (Wed, 17 Jul 2013) $ *)