3. Propositions and Proofs¶
By now, you have seen some ways of defining objects and functions in Lean. In this chapter, we will begin to explain how to write mathematical assertions and proofs in the language of dependent type theory as well.
3.1. Propositions as Types¶
One strategy for proving assertions about objects defined in the language of dependent type theory is to layer an assertion language and a proof language on top of the definition language. But there is no reason to multiply languages in this way: dependent type theory is flexible and expressive, and there is no reason we cannot represent assertions and proofs in the same general framework.
For example, we could introduce a new type, Prop
, to represent propositions, and introduce constructors to build new propositions from others.
constant and : Prop → Prop → Prop
constant or : Prop → Prop → Prop
constant not : Prop → Prop
constant implies : Prop → Prop → Prop
variables p q r : Prop
#check and p q  Prop
#check or (and p q) r  Prop
#check implies (and p q) (and q p)  Prop
We could then introduce, for each element p : Prop
, another type Proof p
, for the type of proofs of p
. An “axiom” would be a constant of such a type.
constant Proof : Prop → Type
constant and_comm : Π p q : Prop,
Proof (implies (and p q) (and q p))
variables p q : Prop
#check and_comm p q  Proof (implies (and p q) (and q p))
In addition to axioms, however, we would also need rules to build new proofs from old ones. For example, in many proof systems for propositional logic, we have the rule of modus ponens:
From a proof ofimplies p q
and a proof ofp
, we obtain a proof ofq
.
We could represent this as follows:
constant modus_ponens :
Π p q : Prop, Proof (implies p q) → Proof p → Proof q
Systems of natural deduction for propositional logic also typically rely on the following rule:
Suppose that, assumingp
as a hypothesis, we have a proof ofq
. Then we can “cancel” the hypothesis and obtain a proof ofimplies p q
.
We could render this as follows:
constant implies_intro :
Π p q : Prop, (Proof p → Proof q) → Proof (implies p q).
This approach would provide us with a reasonable way of building assertions and proofs. Determining that an expression t
is a correct proof of assertion p
would then simply be a matter of checking that t
has type Proof p
.
Some simplifications are possible, however. To start with, we can avoid writing the term Proof
repeatedly by conflating Proof p
with p
itself. In other words, whenever we have p : Prop
, we can interpret p
as a type, namely, the type of its proofs. We can then read t : p
as the assertion that t
is a proof of p
.
Moreover, once we make this identification, the rules for implication show that we can pass back and forth between implies p q
and p → q
. In other words, implication between propositions p
and q
corresponds to having a function that takes any element of p
to an element of q
. As a result, the introduction of the connective implies
is entirely redundant: we can use the usual function space constructor p → q
from dependent type theory as our notion of implication.
This is the approach followed in the Calculus of Constructions, and hence in Lean as well. The fact that the rules for implication in a proof system for natural deduction correspond exactly to the rules governing abstraction and application for functions is an instance of the CurryHoward isomorphism, sometimes known as the propositionsastypes paradigm. In fact, the type Prop
is syntactic sugar for Sort 0
, the very bottom of the type hierarchy described in the last chapter. Moreover, Type u
is also just syntactic sugar for Sort (u+1)
. Prop
has some special features, but like the other type universes, it is closed under the arrow constructor: if we have p q : Prop
, then p → q : Prop
.
There are at least two ways of thinking about propositions as types. To some who take a constructive view of logic and mathematics, this is a faithful rendering of what it means to be a proposition: a proposition p
represents a sort of data type, namely, a specification of the type of data that constitutes a proof. A proof of p
is then simply an object t : p
of the right type.
Those not inclined to this ideology can view it, rather, as a simple coding trick. To each proposition p
we associate a type that is empty if p
is false and has a single element, say *
, if p
is true. In the latter case, let us say that (the type associated with) p
is inhabited. It just so happens that the rules for function application and abstraction can conveniently help us keep track of which elements of Prop
are inhabited. So constructing an element t : p
tells us that p
is indeed true. You can think of the inhabitant of p
as being the “fact that p
is true.” A proof of p → q
uses “the fact that p
is true” to obtain “the fact that q
is true.”
Indeed, if p : Prop
is any proposition, Lean’s kernel treats any two elements t1 t2 : p
as being definitionally equal, much the same way as it treats (λ x, t)s
and t[s/x]
as definitionally equal. This is known as proof irrelevance, and is consistent with the interpretation in the last paragraph. It means that even though we can treat proofs t : p
as ordinary objects in the language of dependent type theory, they carry no information beyond the fact that p
is true.
The two ways we have suggested thinking about the propositionsastypes paradigm differ in a fundamental way. From the constructive point of view, proofs are abstract mathematical objects that are denoted by suitable expressions in dependent type theory. In contrast, if we think in terms of the coding trick described above, then the expressions themselves do not denote anything interesting. Rather, it is the fact that we can write them down and check that they are welltyped that ensures that the proposition in question is true. In other words, the expressions themselves are the proofs.
In the exposition below, we will slip back and forth between these two ways of talking, at times saying that an expression “constructs” or “produces” or “returns” a proof of a proposition, and at other times simply saying that it “is” such a proof. This is similar to the way that computer scientists occasionally blur the distinction between syntax and semantics by saying, at times, that a program “computes” a certain function, and at other times speaking as though the program “is” the function in question.
In any case, all that really matters is the bottom line. To formally express a mathematical assertion in the language of dependent type theory, we need to exhibit a term p : Prop
. To prove that assertion, we need to exhibit a term t : p
. Lean’s task, as a proof assistant, is to help us to construct such a term, t
, and to verify that it is wellformed and has the correct type.
3.2. Working with Propositions as Types¶
In the propositionsastypes paradigm, theorems involving only →
can be proved using lambda abstraction and application. In Lean, the theorem
command introduces a new theorem:
constants p q : Prop
theorem t1 : p → q → p := λ hp : p, λ hq : q, hp
This looks exactly like the definition of the constant function in the last chapter, the only difference being that the arguments are elements of Prop
rather than Type
. Intuitively, our proof of p → q → p
assumes p
and q
are true, and uses the first hypothesis (trivially) to establish that the conclusion, p
, is true.
Note that the theorem
command is really a version of the definition
command: under the propositions and types correspondence, proving the theorem p → q → p
is really the same as defining an element of the associated type. To the kernel type checker, there is no difference between the two.
There are a few pragmatic differences between definitions and theorems, however. In normal circumstances, it is never necessary to unfold the “definition” of a theorem; by proof irrelevance, any two proofs of that theorem are definitionally equal. Once the proof of a theorem is complete, typically we only need to know that the proof exists; it doesn’t matter what the proof is. In light of that fact, Lean tags proofs as irreducible, which serves as a hint to the parser (more precisely, the elaborator) that there is generally no need to unfold it when processing a file. In fact, Lean is generally able to process and check proofs in parallel, since assessing the correctness of one proof does not require knowing the details of another.
As with definitions, the #print
command will show you the proof of a theorem.
theorem t1 : p → q → p := λ hp : p, λ hq : q, hp
#print t1
Notice that the lambda abstractions hp : p
and hq : q
can be viewed as temporary assumptions in the proof of t1
. Lean provides the alternative syntax assume
for such a lambda abstraction:
theorem t1 : p → q → p :=
assume hp : p,
assume hq : q,
hp
Lean also allows us to specify the type of the final term hp
, explicitly, with a show
statement.
theorem t1 : p → q → p :=
assume hp : p,
assume hq : q,
show p, from hp
Adding such extra information can improve the clarity of a proof and help detect errors when writing a proof. The show
command does nothing more than annotate the type, and, internally, all the presentations of t1
that we have seen produce the same term. Lean also allows you to use the alternative syntax lemma
instead of theorem:
lemma t1 : p → q → p :=
assume hp : p,
assume hq : q,
show p, from hp
As with ordinary definitions, we can move the lambdaabstracted variables to the left of the colon:
theorem t1 (hp : p) (hq : q) : p := hp
#check t1  p → q → p
Now we can apply the theorem t1
just as a function application.
axiom hp : p
theorem t2 : q → p := t1 hp
Here, the axiom
command is alternative syntax for constant
. Declaring a “constant” hp : p
is tantamount to declaring that p
is true, as witnessed by hp
. Applying the theorem t1 : p → q → p
to the fact hp : p
that p
is true yields the theorem t2 : q → p
.
Notice, by the way, that the original theorem t1
is true for any propositions p
and q
, not just the particular constants declared. So it would be more natural to define the theorem so that it quantifies over those, too:
theorem t1 (p q : Prop) (hp : p) (hq : q) : p := hp
#check t1
The type of t1
is now ∀ p q : Prop, p → q → p
. We can read this as the assertion “for every pair of propositions p q
, we have p → q → p
.” The symbol ∀
is alternate syntax for Π
, and later we will see how Pi types let us model universal quantifiers more generally. For example, we can move all parameters to the right of the colon:
theorem t1 : ∀ (p q : Prop), p → q → p :=
λ (p q : Prop) (hp : p) (hq : q), hp
If p
and q
have been declared as variables, Lean will generalize them for us automatically:
variables p q : Prop
theorem t1 : p → q → p := λ (hp : p) (hq : q), hp
In fact, by the propositionsastypes correspondence, we can declare the assumption hp
that p
holds, as another variable:
variables p q : Prop
variable hp : p
theorem t1 : q → p := λ (hq : q), hp
Lean detects that the proof uses hp
and automatically adds hp : p
as a premise. In all cases, the command #check t1
still yields ∀ p q : Prop, p → q → p
. Remember that this type can just as well be written ∀ (p q : Prop) (hp : p) (hq :q), p
, since the arrow denotes nothing more than a Pi type in which the target does not depend on the bound variable.
When we generalize t1
in such a way, we can then apply it to different pairs of propositions, to obtain different instances of the general theorem.
theorem t1 (p q : Prop) (hp : p) (hq : q) : p := hp
variables p q r s : Prop
#check t1 p q  p → q → p
#check t1 r s  r → s → r
#check t1 (r → s) (s → r)  (r → s) → (s → r) → r → s
variable h : r → s
#check t1 (r → s) (s → r) h  (s → r) → r → s
Once again, using the propositionsastypes correspondence, the variable h
of type r → s
can be viewed as the hypothesis, or premise, that r → s
holds.
As another example, let us consider the composition function discussed in the last chapter, now with propositions instead of types.
variables p q r s : Prop
theorem t2 (h₁ : q → r) (h₂ : p → q) : p → r :=
assume h₃ : p,
show r, from h₁ (h₂ h₃)
As a theorem of propositional logic, what does t2
say?
Note that it is often useful to use numeric unicode subscripts, entered as \0
, \1
, \2
, …, for hypotheses, as we did in this example.
3.3. Propositional Logic¶
Lean defines all the standard logical connectives and notation. The propositional connectives come with the following notation:
Ascii  Unicode  Editor shortcut  Definition 
true  true  
false  false  
not  ¬  \not , \neg 
not 
/\  ∧  \and 
and 
\/  ∨  \or 
or 
>  →  \to , \r , \imp 

<>  ↔  \iff , \lr 
iff 
They all take values in Prop
.
variables p q : Prop
#check p → q → p ∧ q
#check ¬p → p ↔ false
#check p ∨ q → q ∨ p
The order of operations is as follows: unary negation ¬
binds most strongly, then ∧
, then ∨
, then →
, and finally ↔
. For example, a ∧ b → c ∨ d ∧ e
means (a ∧ b) → (c ∨ (d ∧ e))
. Remember that →
associates to the right (nothing changes now that the arguments are elements of Prop
, instead of some other Type
), as do the other binary connectives. So if we have p q r : Prop
, the expression p → q → r
reads “if p
, then if q
, then r
.” This is just the “curried” form of p ∧ q → r
.
In the last chapter we observed that lambda abstraction can be viewed as an “introduction rule” for →
. In the current setting, it shows how to “introduce” or establish an implication. Application can be viewed as an “elimination rule,” showing how to “eliminate” or use an implication in a proof. The other propositional connectives are defined in Lean’s library in the file init.core
(see Section 6.1 for more information on the library hierarchy), and each connective comes with its canonical introduction and elimination rules.
3.3.1. Conjunction¶
The expression and.intro h1 h2
builds a proof of p ∧ q
using proofs h1 : p
and h2 : q
. It is common to describe and.intro
as the andintroduction rule. In the next example we use and.intro
to create a proof of p → q → p ∧ q
.
example (hp : p) (hq : q) : p ∧ q := and.intro hp hq
#check assume (hp : p) (hq : q), and.intro hp hq
The example
command states a theorem without naming it or storing it in the permanent context. Essentially, it just checks that the given term has the indicated type. It is convenient for illustration, and we will use it often.
The expression and.elim_left h
creates a proof of p
from a proof h : p ∧ q
. Similarly, and.elim_right h
is a proof of q
. They are commonly known as the right and left andelimination rules.
example (h : p ∧ q) : p := and.elim_left h
example (h : p ∧ q) : q := and.elim_right h
Because they are so commonly used, the standard library provides the abbreviations and.left
and and.right
for and.elim_left
and and.elim_right
, respectively.
We can now prove p ∧ q → q ∧ p
with the following proof term.
example (h : p ∧ q) : q ∧ p :=
and.intro (and.right h) (and.left h)
Notice that andintroduction and andelimination are similar to the pairing and projection operations for the cartesian product. The difference is that given hp : p
and hq : q
, and.intro hp hq
has type p ∧ q : Prop
, while pair hp hq
has type p × q : Type
. The similarity between ∧
and ×
is another instance of the CurryHoward isomorphism, but in contrast to implication and the function space constructor, ∧
and ×
are treated separately in Lean. With the analogy, however, the proof we have just constructed is similar to a function that swaps the elements of a pair.
We will see in Chapter 9 that certain types in Lean are structures, which is to say, the type is defined with a single canonical constructor which builds an element of the type from a sequence of suitable arguments. For every p q : Prop
, p ∧ q
is an example: the canonical way to construct an element is to apply and.intro
to suitable arguments hp : p
and hq : q
. Lean allows us to use anonymous constructor notation ⟨arg1, arg2, ...⟩
in situations like these, when the relevant type is an inductive type and can be inferred from the context. In particular, we can often write ⟨hp, hq⟩
instead of and.intro hp hq
:
variables p q : Prop
variables (hp : p) (hq : q)
#check (⟨hp, hq⟩ : p ∧ q)
These angle brackets are obtained by typing \<
and \>
, respectively. Alternatively, you can use ASCII equivalents (
and )
:
variables p q : Prop
variables (hp : p) (hq : q)
example : p ∧ q := (hp, hq)
Lean provides another useful syntactic gadget. Given an expression e
of an inductive type foo
(possibly applied to some arguments), the notation e.bar
is shorthand for foo.bar e
. This provides a convenient way of accessing functions without opening a namespace. For example, the following two expressions mean the same thing:
variable l : list ℕ
#check list.head l
#check l.head
As a result, given h : p ∧ q
, we can write h.left
for and.left h
and h.right
for and.right h
. We can therefore rewrite the sample proof above conveniently as follows:
example (h : p ∧ q) : q ∧ p :=
⟨h.right, h.left⟩
There is a fine line between brevity and obfuscation, and omitting information in this way can sometimes make a proof harder to read. But for straightforward constructions like the one above, when the type of h
and the goal of the construction are salient, the notation is clean and effective.
It is common to iterate constructions like “and.” Lean also allows you to flatten nested constructors that associate to the right, so that these two proofs are equivalent:
example (h : p ∧ q) : q ∧ p ∧ q:=
⟨h.right, ⟨h.left, h.right⟩⟩
example (h : p ∧ q) : q ∧ p ∧ q:=
⟨h.right, h.left, h.right⟩
This is often useful as well.
3.3.2. Disjunction¶
The expression or.intro_left q hp
creates a proof of p ∨ q
from a proof hp : p
. Similarly, or.intro_right p hq
creates a proof for p ∨ q
using a proof hq : q
. These are the left and right orintroduction rules.
example (hp : p) : p ∨ q := or.intro_left q hp
example (hq : q) : p ∨ q := or.intro_right p hq
The orelimination rule is slightly more complicated. The idea is that we can prove r
from p ∨ q
, by showing that r
follows from p
and that r
follows from q
. In other words, it is a proof by cases. In the expression or.elim hpq hpr hqr
, or.elim
takes three arguments, hpq : p ∨ q
, hpr : p → r
and hqr : q → r
, and produces a proof of r
. In the following example, we use or.elim
to prove p ∨ q → q ∨ p
.
example (h : p ∨ q) : q ∨ p :=
or.elim h
(assume hp : p,
show q ∨ p, from or.intro_right q hp)
(assume hq : q,
show q ∨ p, from or.intro_left p hq)
In most cases, the first argument of or.intro_right
and or.intro_left
can be inferred automatically by Lean. Lean therefore provides or.inr
and or.inl
as shorthand for or.intro_right _
and or.intro_left _
. Thus the proof term above could be written more concisely:
example (h : p ∨ q) : q ∨ p :=
or.elim h (λ hp, or.inr hp) (λ hq, or.inl hq)
Notice that there is enough information in the full expression for Lean to infer the types of hp
and hq
as well. But using the type annotations in the longer version makes the proof more readable, and can help catch and debug errors.
Because or
has two constructors, we cannot use anonymous constructor notation. But we can still write h.elim
instead of or.elim h
:
example (h : p ∨ q) : q ∨ p :=
h.elim
(assume hp : p, or.inr hp)
(assume hq : q, or.inl hq)
Once again, you should exercise judgment as to whether such abbreviations enhance or diminish readability.
3.3.3. Negation and Falsity¶
Negation, ¬p
, is actually defined to be p → false
, so we obtain ¬p
by deriving a contradiction from p
. Similarly, the expression hnp hp
produces a proof of false
from hp : p
and hnp : ¬p
. The next example uses both these rules to produce a proof of (p → q) → ¬q → ¬p
. (The symbol ¬
is produced by typing \not
or \neg
.)
example (hpq : p → q) (hnq : ¬q) : ¬p :=
assume hp : p,
show false, from hnq (hpq hp)
The connective false
has a single elimination rule, false.elim
, which expresses the fact that anything follows from a contradiction. This rule is sometimes called ex falso (short for ex falso sequitur quodlibet), or the principle of explosion.
example (hp : p) (hnp : ¬p) : q := false.elim (hnp hp)
The arbitrary fact, q
, that follows from falsity is an implicit argument in false.elim
and is inferred automatically. This pattern, deriving an arbitrary fact from contradictory hypotheses, is quite common, and is represented by absurd
.
example (hp : p) (hnp : ¬p) : q := absurd hp hnp
Here, for example, is a proof of ¬p → q → (q → p) → r
:
example (hnp : ¬p) (hq : q) (hqp : q → p) : r :=
absurd (hqp hq) hnp
Incidentally, just as false
has only an elimination rule, true
has only an introduction rule, true.intro : true
, sometimes abbreviated trivial : true
. In other words, true
is simply true, and has a canonical proof, trivial
.
3.3.4. Logical Equivalence¶
The expression iff.intro h1 h2
produces a proof of p ↔ q
from h1 : p → q
and h2 : q → p
. The expression iff.elim_left h
produces a proof of p → q
from h : p ↔ q
. Similarly, iff.elim_right h
produces a proof of q → p
from h : p ↔ q
. Here is a proof of p ∧ q ↔ q ∧ p
:
theorem and_swap : p ∧ q ↔ q ∧ p :=
iff.intro
(assume h : p ∧ q,
show q ∧ p, from and.intro (and.right h) (and.left h))
(assume h : q ∧ p,
show p ∧ q, from and.intro (and.right h) (and.left h))
#check and_swap p q  p ∧ q ↔ q ∧ p
Because they represent a form of modus ponens, iff.elim_left
and iff.elim_right
can be abbreviated iff.mp
and iff.mpr
, respectively. In the next example, we use that theorem to derive q ∧ p
from p ∧ q
:
variable h : p ∧ q
example : q ∧ p := iff.mp (and_swap p q) h
We can use the anonymous constructor notation to construct a proof of p ↔ q
from proofs of the forward and backward directions, and we can also use .
notation with mp
and mpr
. The previous examples can therefore be written concisely as follows:
theorem and_swap : p ∧ q ↔ q ∧ p :=
⟨ λ h, ⟨h.right, h.left⟩, λ h, ⟨h.right, h.left⟩ ⟩
example (h : p ∧ q) : q ∧ p := (and_swap p q).mp h
3.4. Introducing Auxiliary Subgoals¶
This is a good place to introduce another device Lean offers to help structure long proofs, namely, the have
construct, which introduces an auxiliary subgoal in a proof. Here is a small example, adapted from the last section:
variables p q : Prop
example (h : p ∧ q) : q ∧ p :=
have hp : p, from and.left h,
have hq : q, from and.right h,
show q ∧ p, from and.intro hq hp
Internally, the expression have h : p, from s, t
produces the term (λ (h : p), t) s
. In other words, s
is a proof of p
, t
is a proof of the desired conclusion assuming h : p
, and the two are combined by a lambda abstraction and application. This simple device is extremely useful when it comes to structuring long proofs, since we can use intermediate have
’s as stepping stones leading to the final goal.
Lean also supports a structured way of reasoning backwards from a goal, which models the “suffices to show” construction in ordinary mathematics. The next example simply permutes the last two lines in the previous proof.
variables p q : Prop
example (h : p ∧ q) : q ∧ p :=
have hp : p, from and.left h,
suffices hq : q, from and.intro hq hp,
show q, from and.right h
Writing suffices hq : q
leaves us with two goals. First, we have to show that it indeed suffices to show q
, by proving the original goal of q ∧ p
with the additional hypothesis hq : q
. Finally, we have to show q
.
3.5. Classical Logic¶
The introduction and elimination rules we have seen so far are all constructive, which is to say, they reflect a computational understanding of the logical connectives based on the propositionsastypes correspondence. Ordinary classical logic adds to this the law of the excluded middle, p ∨ ¬p
. To use this principle, you have to open the classical namespace.
open classical
variable p : Prop
#check em p
Intuitively, the constructive “or” is very strong: asserting p ∨ q
amounts to knowing which is the case. If RH
represents the Riemann hypothesis, a classical mathematician is willing to assert RH ∨ ¬RH
, even though we cannot yet assert either disjunct.
One consequence of the law of the excluded middle is the principle of doublenegation elimination:
theorem dne {p : Prop} (h : ¬¬p) : p :=
or.elim (em p)
(assume hp : p, hp)
(assume hnp : ¬p, absurd hnp h)
Doublenegation elimination allows one to prove any proposition, p
, by assuming ¬p
and deriving false
, because that amounts to proving ¬¬p
. In other words, doublenegation elimination allows one to carry out a proof by contradiction, something which is not generally possible in constructive logic. As an exercise, you might try proving the converse, that is, showing that em
can be proved from dne
.
The classical axioms also give you access to additional patterns of proof that can be justified by appeal to em
. For example, one can carry out a proof by cases:
example (h : ¬¬p) : p :=
by_cases
(assume h1 : p, h1)
(assume h1 : ¬p, absurd h1 h)
Or you can carry out a proof by contradiction:
example (h : ¬¬p) : p :=
by_contradiction
(assume h1 : ¬p,
show false, from h h1)
If you are not used to thinking constructively, it may take some time for you to get a sense of where classical reasoning is used. It is needed in the following example because, from a constructive standpoint, knowing that p
and q
are not both true does not necessarily tell you which one is false:
example (h : ¬(p ∧ q)) : ¬p ∨ ¬q :=
or.elim (em p)
(assume hp : p,
or.inr
(show ¬q, from
assume hq : q,
h ⟨hp, hq⟩))
(assume hp : ¬p,
or.inl hp)
We will see later that there are situations in constructive logic where principles like excluded middle and doublenegation elimination are permissible, and Lean supports the use of classical reasoning in such contexts without relying on excluded middle.
The full list of axioms that are used in Lean to support classical reasoning are discussed in Chapter 11.
3.6. Examples of Propositional Validities¶
Lean’s standard library contains proofs of many valid statements of propositional logic, all of which you are free to use in proofs of your own. The following list includes a number of common identities. The ones that require classical reasoning are grouped together at the end, while the rest are constructively valid.
open classical
variables p q r s : Prop
 commutativity of ∧ and ∨
example : p ∧ q ↔ q ∧ p := sorry
example : p ∨ q ↔ q ∨ p := sorry
 associativity of ∧ and ∨
example : (p ∧ q) ∧ r ↔ p ∧ (q ∧ r) := sorry
example : (p ∨ q) ∨ r ↔ p ∨ (q ∨ r) := sorry
 distributivity
example : p ∧ (q ∨ r) ↔ (p ∧ q) ∨ (p ∧ r) := sorry
example : p ∨ (q ∧ r) ↔ (p ∨ q) ∧ (p ∨ r) := sorry
 other properties
example : (p → (q → r)) ↔ (p ∧ q → r) := sorry
example : ((p ∨ q) → r) ↔ (p → r) ∧ (q → r) := sorry
example : ¬(p ∨ q) ↔ ¬p ∧ ¬q := sorry
example : ¬p ∨ ¬q → ¬(p ∧ q) := sorry
example : ¬(p ∧ ¬p) := sorry
example : p ∧ ¬q → ¬(p → q) := sorry
example : ¬p → (p → q) := sorry
example : (¬p ∨ q) → (p → q) := sorry
example : p ∨ false ↔ p := sorry
example : p ∧ false ↔ false := sorry
example : ¬(p ↔ ¬p) := sorry
example : (p → q) → (¬q → ¬p) := sorry
 these require classical reasoning
example : (p → r ∨ s) → ((p → r) ∨ (p → s)) := sorry
example : ¬(p ∧ q) → ¬p ∨ ¬q := sorry
example : ¬(p → q) → p ∧ ¬q := sorry
example : (p → q) → (¬p ∨ q) := sorry
example : (¬q → ¬p) → (p → q) := sorry
example : p ∨ ¬p := sorry
example : (((p → q) → p) → p) := sorry
The sorry
identifier magically produces a proof of anything, or provides an object of any data type at all. Of course, it is unsound as a proof method – for example, you can use it to prove false
– and Lean produces severe warnings when files use or import theorems which depend on it. But it is very useful for building long proofs incrementally. Start writing the proof from the top down, using sorry
to fill in subproofs. Make sure Lean accepts the term with all the sorry
’s; if not, there are errors that you need to correct. Then go back and replace each sorry
with an actual proof, until no more remain.
Here is another useful trick. Instead of using sorry
, you can use an underscore _
as a placeholder. Recall that this tells Lean that the argument is implicit, and should be filled in automatically. If Lean tries to do so and fails, it returns with an error message “don’t know how to synthesize placeholder.” This is followed by the type of the term it is expecting, and all the objects and hypothesis available in the context. In other words, for each unresolved placeholder, Lean reports the subgoal that needs to be filled at that point. You can then construct a proof by incrementally filling in these placeholders.
For reference, here are two sample proofs of validities taken from the list above.
open classical
variables p q r : Prop
 distributivity
example : p ∧ (q ∨ r) ↔ (p ∧ q) ∨ (p ∧ r) :=
iff.intro
(assume h : p ∧ (q ∨ r),
have hp : p, from h.left,
or.elim (h.right)
(assume hq : q,
show (p ∧ q) ∨ (p ∧ r), from or.inl ⟨hp, hq⟩)
(assume hr : r,
show (p ∧ q) ∨ (p ∧ r), from or.inr ⟨hp, hr⟩))
(assume h : (p ∧ q) ∨ (p ∧ r),
or.elim h
(assume hpq : p ∧ q,
have hp : p, from hpq.left,
have hq : q, from hpq.right,
show p ∧ (q ∨ r), from ⟨hp, or.inl hq⟩)
(assume hpr : p ∧ r,
have hp : p, from hpr.left,
have hr : r, from hpr.right,
show p ∧ (q ∨ r), from ⟨hp, or.inr hr⟩))
 an example that requires classical reasoning
example : ¬(p ∧ ¬q) → (p → q) :=
assume h : ¬(p ∧ ¬q),
assume hp : p,
show q, from
or.elim (em q)
(assume hq : q, hq)
(assume hnq : ¬q, absurd (and.intro hp hnq) h)
3.7. Exercises¶
 Prove as many identities from the previous section as you can, replacing the “sorry” placeholders with actual proofs.
 Prove
¬(p ↔ ¬p)
without using classical logic.