Volume 1: Logical Foundations

# IndPrinciplesInduction Principles

Every time we declare a new Inductive datatype, Coq automatically generates an induction principle for this type. This induction principle is a theorem like any other: If t is defined inductively, the corresponding induction principle is called t_ind.

# Basics

Here is the induction principle for natural numbers:
Check nat_ind :
∀ P : nat → Prop,
P 0 →
(∀ n : nat, P n → P (S n)) →
∀ n : nat, P n.
In English: Suppose P is a property of natural numbers (that is, P n is a Prop for every n). To show that P n holds of all n, it suffices to show:

- P holds of 0
- for any n, if P holds of n, then P holds of S 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 with directly using apply nat_ind, 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 Induction chapter.
Theorem mul_0_r' : ∀ n:nat,
n × 0 = 0.
Proof.
apply nat_ind.
- (* n = O *) reflexivity.
- (* n = S n' *) 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. By contrast, the induction tactic works either with a variable in the context or a quantified variable in the goal.
Third, we had to manually supply the name of the induction principle with apply, but induction figures that out itself.
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 these bits of bookkeeping, applying nat_ind is what we are really doing.

#### Exercise: 2 stars, standard (plus_one_r')

Complete this proof 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 of course we don't need the proof technique of 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:
Inductive time : Type :=
| day
| night.

Check time_ind :
∀ P : time → Prop,
P day →
P night →
∀ t : time, P t.

#### Exercise: 1 star, standard, 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
| green
| blue.
Check rgb_ind.
☐
Here's another example, this time with one of the constructors taking some arguments.
Inductive natlist : Type :=
| nnil
| ncons (n : nat) (l : natlist).

Check natlist_ind :
∀ P : natlist → Prop,
P nnil →
(∀ (n : nat) (l : natlist),
P l → P (ncons n l)) →
∀ l : natlist, P l.
In general, the automatically generated induction principle for inductive type t is formed as follows:

- Each constructor c generates one case of the principle.
- If c takes no arguments, that case is:
  "P holds of c"
- If c takes arguments x1:a1 ... xn:an, that case is:
  "For all x1:a1 ... xn:an, if [P] holds of each of the arguments of type [t], then [P] holds of [c x1 ... xn]" But that oversimplifies a little. An assumption about P holding of an argument x of type t actually occurs immediately after the quantification of x.

For example, suppose we had written the definition of natlist a little differently:
Inductive natlist' : Type :=
| nnil'
| nsnoc (l : natlist') (n : nat).
Now the induction principle case for nsnoc is a bit different than the earlier case for ncons:
Check natlist'_ind :
∀ P : natlist' → Prop,
P nnil' →
(∀ l : natlist', P l → ∀ n : nat, P (nsnoc l n)) →
∀ n : natlist', P n.

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

Here is a type for trees that contain a boolean value at each leaf and branch.
Inductive booltree : Type :=
| bt_empty
| bt_leaf (b : bool)
| bt_branch (b : bool) (t1 t2 : booltree).

(* What is the induction principle for booltree? Of course you could
ask Coq, but try not to do that. Instead, write it down yourself on
paper. Then look at the definition of booltree_ind_type, below.
It has three missing pieces, which are provided by the definitions
in between here and there. Fill in those definitions based on what
you wrote on paper. *)

Definition booltree_property_type : Type := booltree → Prop.

Definition base_case (P : booltree_property_type) : Prop
(* REPLACE THIS LINE WITH ":= _your_definition_ ." *). Admitted.

Definition leaf_case (P : booltree_property_type) : Prop
(* REPLACE THIS LINE WITH ":= _your_definition_ ." *). Admitted.

Definition branch_case (P : booltree_property_type) : Prop
(* REPLACE THIS LINE WITH ":= _your_definition_ ." *). Admitted.

Definition booltree_ind_type :=
∀ (P : booltree_property_type),
base_case P →
leaf_case P →
branch_case P →
∀ (b : booltree), P b.
Now check the correctness of your answers by proving the following theorem. If you have them right, you can complete the proof with just one tactic: exact booltree_ind. That will work because the automatically generated induction principle booltree_ind has the same type as what you just defined.
Theorem booltree_ind_type_correct : booltree_ind_type.
Proof. (* FILL IN HERE *) Admitted.
☐

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

Here is an induction principle for a toy type:
∀ P : Toy → Prop,
(∀ b : bool, P (con1 b)) →
(∀ (n : nat) (t : Toy), P t → P (con2 n t)) →
∀ t : Toy, P t Give an Inductive definition of Toy, such that the induction principle Coq generates is that given above:
Inductive Toy : Type :=
(* FILL IN HERE *)
.
Show that your definition is correct by proving the following theorem. You should be able to instantiate f and g with your two constructors, then immediately finish the proof with exact Toy_ind. As in the previous exercise, that will work because the automatically generated induction principle Toy_ind will have the same type.
Theorem Toy_correct : ∃ f g,
∀ P : Toy → Prop,
(∀ b : bool, P (f b)) →
(∀ (n : nat) (t : Toy), P t → P (g n t)) →
∀ t : Toy, P t.
Proof. (* FILL IN HERE *) Admitted.
☐

# Polymorphism

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:
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 that 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.

#### Exercise: 1 star, standard, 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 : X)
| node (t1 t2 : tree X).
Check tree_ind.
☐

#### Exercise: 1 star, standard, 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 ☐

#### Exercise: 1 star, standard, 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 ☐

#### Exercise: 1 star, standard, optional (foo')

Consider the following inductive definition:
Inductive foo' (X:Type) : Type :=
| C1 (l : list X) (f : foo' X)
| C2.
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, ________________________
☐

# Induction Hypotheses

Where does the phrase "induction hypothesis" fit into this story?
The induction principle for numbers
∀ 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 (or rather, 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.
We can make proofs by induction more explicit by giving this expression a name. For example, instead of stating the theorem mul_0_r as "∀ n, n × 0 = 0," we can write it as "∀ n, P_m0r n", where P_m0r is defined as...
Definition P_m0r (n:nat) : Prop :=
n × 0 = 0.
... or equivalently:
Definition P_m0r' : nat → Prop :=
fun n ⇒ n × 0 = 0.
Now it is easier to see where P_m0r appears in the proof.
Theorem mul_0_r'' : ∀ n:nat,
P_m0r n.
Proof.
apply nat_ind.
- (* n = O *) reflexivity.
- (* 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 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 (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

The induction tactic actually does even more low-level bookkeeping for us than we discussed above.
Recall the informal statement of the induction principle for natural numbers:

- 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.

So, when we begin a proof with intros n and then induction n, we are first telling Coq to consider a particular n (by introducing it into the context) and then telling it to prove something about all numbers (by using induction).
What Coq actually does in this situation, internally, is it "re-generalizes" the variable we perform induction on. For example, in our original proof that plus is associative...
Theorem add_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'].
- (* n = O *) reflexivity.
- (* n = S n' *)
simpl. rewrite → IHn'. reflexivity. Qed.
It also works to apply induction to a variable that is quantified in the goal.
Theorem add_comm' : ∀ n m : nat,
n + m = m + n.
Proof.
induction n as [| n'].
- (* n = O *) intros m. rewrite → add_0_r. reflexivity.
- (* 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 add_comm'' : ∀ n m : nat,
n + m = m + n.
Proof.
(* Let's do induction on m this time, instead of n... *)
induction m as [| m']. (* n is already introduced into the context *)
- (* m = O *) simpl. rewrite → add_0_r. reflexivity.
- (* m = S m' *) simpl. rewrite <- IHm'.
rewrite <- plus_n_Sm. reflexivity. Qed.

#### Exercise: 1 star, standard, optional (plus_explicit_prop)

Rewrite both add_assoc' and add_comm' and their proofs in the same style as mul_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 *)
☐

# Induction Principles for Propositions

Inductive definitions of propositions also cause Coq to generate induction priniciples. For example, recall our proposition ev from IndProp:
Print ev.

(* ===>
Inductive ev : nat -> Prop :=
| ev_0 : ev 0
| ev_SS : forall n : nat, ev n -> ev (S (S n)))
*)

Check ev_ind :
∀ P : nat → Prop,
P 0 →
(∀ n : nat, ev n → P n → P (S (S n))) →
∀ n : nat, ev n → P n.
In English, ev_ind says: Suppose P is a property of natural numbers. To show that P n holds whenever n is even, it suffices to show:

- P holds for 0,
- for any n, if n is even and P holds for n, then P holds for S (S n).

As expected, we can apply ev_ind directly instead of using induction. For example, we can use it to show that ev' (the slightly awkward alternate definition of evenness that we saw in an exercise in the IndProp chapter) is equivalent to the cleaner inductive definition ev:
Inductive ev' : nat → Prop :=
| ev'_0 : ev' 0
| ev'_2 : ev' 2
| ev'_sum n m (Hn : ev' n) (Hm : ev' m) : ev' (n + m).

Theorem ev_ev' : ∀ n, ev n → ev' n.
Proof.
apply ev_ind.
- (* ev_0 *)
apply ev'_0.
- (* ev_SS *)
intros m Hm IH.
apply (ev'_sum 2 m).
+ apply ev'_2.
+ apply IH.
Qed.
The precise form of an Inductive definition can affect the induction principle Coq generates.
Inductive le1 : nat → nat → Prop :=
| le1_n : ∀ n, le1 n n
| le1_S : ∀ n m, (le1 n m) → (le1 n (S m)).

Notation "m <=1 n" := (le1 m n) (at level 70).
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 le2 (n:nat) : nat → Prop :=
| le2_n : le2 n n
| le2_S m (H : le2 n m) : le2 n (S m).

Notation "m <=2 n" := (le2 m n) (at level 70).
The second one is better, even though it looks less symmetric. Why? Because it gives us a simpler induction principle.
Check le1_ind :
∀ P : nat → nat → Prop,
(∀ n : nat, P n n) →
(∀ n m : nat, n <=1 m → P n m → P n (S m)) →
∀ n n0 : nat, n <=1 n0 → P n n0.

Check le2_ind :
∀ (n : nat) (P : nat → Prop),
P n →
(∀ m : nat, n <=2 m → P m → P (S m)) →
∀ n0 : nat, n <=2 n0 → P n0.

# Another Form of Induction Principles on Propositions (Optional)

The induction principle that Coq generated for ev was parameterized on a natural number n. It could have additionally been parameterized on the evidence that n was even, which would have led to this induction principle:
∀ P : (∀ n : nat, ev'' n → Prop),
P O ev_0 →
(∀ (m : nat) (E : ev'' m),
P m E → P (S (S m)) (ev_SS m E)) →
∀ (n : nat) (E : ev'' n), P n E
... because:

- Since ev is indexed by a number n (every ev object E is a piece of evidence that some particular number n is even), the proposition P is parameterized by both n and E -- that is, the induction principle can be used to prove assertions involving both an even number and the evidence that it is even.
- Since there are two ways of giving evidence of evenness (even has two constructors), applying the induction principle generates two subgoals:
  - We must prove that P holds for O and ev_0.
  - We must prove that, whenever m is an even number and E is an evidence of its evenness, if P holds of m and E, then it also holds of S (S m) and ev_SS m E.
- If these subgoals can be proved, then the induction principle tells us that P is true for all even numbers n and evidence E of their evenness.

This is more flexibility than we normally need or want: it is giving us a way to prove logical assertions where the assertion involves properties of some piece of evidence of evenness, while all we really care about is proving properties of numbers that are even -- we are interested in assertions about numbers, not about evidence. It would therefore be more convenient to have an induction principle for proving propositions P that are parameterized just by n and whose conclusion establishes P for all even numbers n:
∀ P : nat → Prop,
... →
∀ n : nat,
even n → P n That is why Coq actually generates the induction principle ev_ind that we saw before.

# Formal vs. Informal Proofs by Induction

Question: What is the relation between a formal proof of a proposition P and an informal proof of the same proposition P?
Answer: The latter should teach the reader everything they would need to understand to be able to produce the former.
Question: How much detail does that require?
Unfortunately, there is no single right answer; rather, there is a range of choices.
At one end of the spectrum, we can essentially give the reader the whole formal proof (i.e., the "informal" proof will amount to just transcribing the formal one into words). This may give the reader the ability to reproduce the formal one for themselves, but it probably doesn't teach them anything much.
At the other end of the spectrum, we can say "The theorem is true and you can figure out why for yourself if you think about it hard enough." This is also not a good teaching strategy, because often writing the proof requires one or more significant insights into the thing we're proving, and most readers will give up before they rediscover all the same insights as we did.
In the middle is the golden mean -- a proof that includes all of the essential insights (saving the reader the hard work that we went through to find the proof in the first place) plus high-level suggestions for the more routine parts to save the reader from spending too much time reconstructing these (e.g., what the IH says and what must be shown in each case of an inductive proof), but not so much detail that the main ideas are obscured.
Since we've spent much of this chapter looking "under the hood" at formal proofs by induction, now is a good moment to talk a little about informal proofs by induction.
In the real world of mathematical communication, written proofs range from extremely longwinded and pedantic to extremely brief and telegraphic. Although the ideal is somewhere in between, while one is getting used to the style it is better to start out at the pedantic end. Also, during the learning phase, it is probably helpful to have a clear standard to compare against. With this in mind, we offer two templates -- one for proofs by induction over data (i.e., where the thing we're doing induction on lives in Type) and one for proofs by induction over evidence (i.e., where the inductively defined thing lives in Prop).

## Induction Over an Inductively Defined Set

Template:

- Theorem: <Universally quantified proposition of the form "For all n:S, P(n)," where S is some inductively defined set.>
  Proof: By induction on n.
  <one case for each constructor c of S...>
  - 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)>.
    <go on and prove P(n) to finish the case...>
  - <other cases similarly...> ☐

Example:

- Theorem: For all sets X, lists l : list X, and numbers n, if length l = n then index (S n) l = None.
  Proof: By induction on l.
  - Suppose l = []. We must show, for all numbers n, that, if length [] = n, then index (S n) [] = None.
    This follows immediately from the definition of index.
  - 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.
    Let n be a number with length l = n. Since
    length l = length (x::l') = S (length l'), it suffices to show that
    index (S (length l')) l' = None. But this follows directly from the induction hypothesis, picking n' to be length l'. ☐

## Induction Over an Inductively Defined Proposition

Since inductively defined proof objects are often called "derivation trees," this form of proof is also known as induction on derivations.
Template:

- 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")>
  Proof: By induction on a derivation of Q. <Or, more generally, "Suppose we are given x, y, and z. We show that Q x y z implies P x y z, by induction on a derivation of Q x y z"...>
  <one case for each constructor c of Q...>
  - 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>.
    <go on and prove P to finish the case...>
  - <other cases similarly...> ☐

Example

- Theorem: The ≤ relation is transitive -- i.e., for all numbers n, m, and o, if n ≤ m and m ≤ o, then n ≤ o.
  Proof: By induction on a derivation of m ≤ 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'.
    But then, by le_S, n ≤ S o'. ☐

# Explicit Proof Objects for Induction (Optional)

Although tactic-based proofs are normally much easier to work with, the ability to write a proof term directly is sometimes very handy, particularly when we want Coq to do something slightly non-standard.
Recall again the induction principle on naturals that Coq generates for us automatically from the Inductive declaration for nat.
Check nat_ind :
∀ P : nat → Prop,
P 0 →
(∀ n : nat, P n → P (S n)) →
∀ 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.
We can rewrite that more tidily as follows:
Fixpoint build_proof
(P : nat → Prop)
(evPO : P 0)
(evPS : ∀ n : nat, P n → P (S n))
(n : nat) : P n :=
match n with
| 0 ⇒ evPO
| S k ⇒ evPS k (build_proof P evPO evPS k)
end.

Definition nat_ind_tidy := build_proof.
We can read build_proof as follows: Suppose we have evidence evPO that P holds on 0, and evidence evPS that ∀ n:nat, P n → P (S n). Then we can prove that P holds of an arbitrary nat n using recursive function build_proof, which pattern matches on n:

- If n is 0, build_proof returns evPO to show that P n holds.
- If n is S k, build_proof applies itself recursively on k to obtain evidence that P k holds; then it applies evPS on that evidence to show that P (S n) holds.

Recursive function build_proof thus pattern matches against n, recursing all the way down to 0, and building up a proof as it returns.
The actual nat_ind that Coq generates uses a recursive function F defined with fix instead of Fixpoint.
We can adapt this approach to proving nat_ind to help prove non-standard induction principles too. As a motivating example, suppose that we want to prove the following lemma, directly relating the ev predicate we defined in IndProp to the even function defined in Basics.
Lemma even_ev : ∀ n: nat, even n = true → ev n.
Proof.
induction n; intros.
- apply ev_0.
- destruct n.
+ simpl in H. inversion H.
+ simpl in H.
apply ev_SS.
Abort.
Attempts to prove this by standard induction on n fail in the case for S (S n), because the induction hypothesis only tells us something about S n, which is useless. There are various ways to hack around this problem; for example, we can use ordinary induction on n to prove this (try it!):
Lemma even_ev' : ∀ n : nat, (even n = true → ev n) ∧ (even (S n) = true → ev (S n)).
But we can make a much better proof by defining and proving a non-standard induction principle that goes "by twos":
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.
The induction ... using tactic variant gives a convenient way to utilize a non-standard induction principle like this.
Lemma even_ev : ∀ n, even n = true → ev n.
Proof.
intros.
induction n as [ | |n'] using nat_ind2.
- apply ev_0.
- simpl in H.
inversion H.
- simpl in H.
apply ev_SS.
apply IHn'.
apply H.
Qed.

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

What if we wanted to define binary trees as follows, using a constructor that bundles the children and value at a node into a tuple?
Notation "( x , y , .. , z )" := (pair .. (pair x y) .. z) : core_scope.

Inductive t_tree (X : Type) : Type :=
| t_leaf
| t_branch : (t_tree X × X × t_tree X) → t_tree X.

Arguments t_leaf {X}.
Arguments t_branch {X}.
Unfortunately, the automatically-generated induction principle is not as strong as we need. It doesn't introduce induction hypotheses for the subtrees.
Check t_tree_ind.
That will get us in trouble if we want to prove something by induction, such as that reflect is an involution.
Fixpoint reflect {X : Type} (t : t_tree X) : t_tree X :=
match t with
| t_leaf ⇒ t_leaf
| t_branch (l, v, r) ⇒ t_branch (reflect r, v, reflect l)
end.

Theorem reflect_involution : ∀ (X : Type) (t : t_tree X),
reflect (reflect t) = t.
Proof.
intros X t. induction t.
- reflexivity.
- destruct p as [[l v] r]. simpl. Abort.
We get stuck, because we have no inductive hypothesis for l or r. So, we need to define our own custom induction principle, and use it to complete the proof.
First, define the type of the induction principle that you want to use. There are many possible answers. Recall that you can use match as part of the definition.
Definition better_t_tree_ind_type : Prop
(* REPLACE THIS LINE WITH ":= _your_definition_ ." *). Admitted.
Second, define the induction principle by giving a term of that type. Use the examples about nat, above, as models.
Definition better_t_tree_ind : better_t_tree_ind_type
(* REPLACE THIS LINE WITH ":= _your_definition_ ." *). Admitted.
Finally, prove the theorem. If induction...using gives you an error about "Cannot recognize an induction scheme", don't worry about it. The induction tactic is picky about the shape of the theorem you pass to it, but it doesn't give you much information to debug what is wrong about that shape. You can use apply instead, as we saw at the beginning of this file.
Theorem reflect_involution : ∀ (X : Type) (t : t_tree X),
reflect (reflect t) = t.
Proof. (* FILL IN HERE *) Admitted.
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(* 2023-09-11 13:53 *)