## Caveats

I really don't know much category theory. Just enough to put this
together. Also, this really is "put together." I haven't yet found an
authoritative source (that's accessible to a category theory beginner like
myself) that discusses the correspondence between the category-theoretic and
functional programming uses of these notions in enough detail to be sure that
none of the pieces here is mistaken.
In particular, it wasn't completely obvious how to map the polymorphism on the
programming theory side into the category theory. The way I accomplished this
may be more complex than it needs to be.
Also I'm bothered by the fact that our `<=<`

operation is only partly defined
on our domain of natural transformations.
There are three additional points below that I wonder whether may be too
cavalier.
But all considered, this does seem to
me to be a reasonable way to put the pieces together. We very much welcome
feedback from anyone who understands these issues better, and will make
corrections.

Thanks Wren Thornton for helpful comments on these notes (not yet incorporated).

This page was a helpful starting point.

## Monoids

A **monoid** is a structure `(S,⋆,z)`

consisting of an associative binary operation `⋆`

over some set `S`

, which is closed under `⋆`

, and which contains an identity element `z`

for `⋆`

. That is:

for all s1, s2, s3 in S: (i) s1⋆s2 etc are also in S (ii) (s1⋆s2)⋆s3 = s1⋆(s2⋆s3) (iii) z⋆s1 = s1 = s1⋆z

Some examples of monoids are:

- finite strings of an alphabet
`A`

, with`⋆`

being concatenation and`z`

being the empty string - all functions
`X→X`

over a set`X`

, with`⋆`

being composition and`z`

being the identity function over`X`

- the natural numbers with
`⋆`

being plus and`z`

being 0 (in particular, this is a**commutative monoid**). If we use the integers, or the naturals mod n, instead of the naturals, then every element will have an inverse and so we have not merely a monoid but a**group**. - if we let
`⋆`

be multiplication and`z`

be 1, we get different monoids over the same sets as in the previous item.

## Categories

A **category** is a generalization of a monoid. A category consists of a class of **elements**, and a class of **morphisms** between those elements. Morphisms are sometimes also called maps or arrows. They are something like functions (and as we'll see below, given a set of functions they'll determine a category). However, a single morphism only maps between a single source element and a single target element. Also, there can be multiple distinct morphisms between the same source and target, so the identity of a morphism goes beyond its "extension."

When a morphism `f`

in category **C** has source `C1`

and target `C2`

, we'll write `f:C1→C2`

.

To have a category, the elements and morphisms have to satisfy some constraints:

(i) the class of morphisms has to be closed under composition: where f:C1→C2 and g:C2→C3, g ∘ f is also a morphism of the category, which maps C1→C3. (ii) composition of morphisms has to be associative (iii) every element X of the category has to have an identity morphism 1_{X}, which is such that for every morphism f:C1→C2: 1_{C2}∘ f = f = f ∘ 1_{C1}

These parallel the constraints for monoids. Note that there can be multiple distinct morphisms between an element `X`

and itself; they need not all be identity morphisms. Indeed from (iii) it follows that each element can have only a single identity morphism.

A good intuitive picture of a category is as a generalized directed graph, where the category elements are the graph's nodes, and there can be multiple directed edges between a given pair of nodes, and nodes can also have multiple directed edges to themselves. Morphisms correspond to directed paths of length ≥ 0 in the graph.

Some examples of categories are:

Categories whose elements are sets and whose morphisms are functions between those sets. Here the source and target of a function are its domain and range, so distinct functions sharing a domain and range (e.g.,

`sin`

and`cos`

) are distinct morphisms between the same source and target elements. The identity morphism for any element/set is just the identity function for that set.any monoid

`(S,⋆,z)`

generates a category with a single element`Q`

; this`Q`

need not have any relation to`S`

. The members of`S`

play the role of*morphisms*of this category, rather than its elements. All of these morphisms are understood to map`Q`

to itself. The result of composing the morphism consisting of`s1`

with the morphism`s2`

is the morphism`s3`

, where`s3=s1⋆s2`

. The identity morphism for the (single) category element`Q`

is the monoid's identity`z`

.a

**preorder**is a structure`(S, ≤)`

consisting of a reflexive, transitive, binary relation on a set`S`

. It need not be connected (that is, there may be members`s1`

,`s2`

of`S`

such that neither`s1 ≤ s2`

nor`s2 ≤ s1`

). It need not be anti-symmetric (that is, there may be members`s1`

,`s2`

of`S`

such that`s1 ≤ s2`

and`s2 ≤ s1`

but`s1`

and`s2`

are not identical). Some examples:- sentences ordered by logical implication ("p and p" implies and is implied by "p", but these sentences are not identical; so this illustrates a pre-order without anti-symmetry)
- sets ordered by size (this illustrates it too)

Any pre-order

`(S,≤)`

generates a category whose elements are the members of`S`

and which has only a single morphism between any two elements`s1`

and`s2`

, iff`s1 ≤ s2`

.

## Functors

A **functor** is a "homomorphism", that is, a structure-preserving mapping, between categories. In particular, a functor `F`

from category **C** to category **D** must:

(i) associate with every element C1 ofCan element F(C1) ofD(ii) associate with every morphism f:C1→C2 ofCa morphism F(f):F(C1)→F(C2) ofD(iii) "preserve identity", that is, for every element C1 ofC: F of C1's identity morphism inCmust be the identity morphism of F(C1) inD: F(1_{C1}) = 1_{F(C1)}. (iv) "distribute over composition", that is for any morphisms f and g inC: F(g ∘ f) = F(g) ∘ F(f)

A functor that maps a category to itself is called an **endofunctor**. The (endo)functor that maps every element and morphism of **C** to itself is denoted `1C`

.

How functors compose: If `G`

is a functor from category **C** to category **D**, and `K`

is a functor from category **D** to category **E**, then `KG`

is a functor which maps every element `C1`

of **C** to element `K(G(C1))`

of **E**, and maps every morphism `f`

of **C** to morphism `K(G(f))`

of **E**.

I'll assert without proving that functor composition is associative.

## Natural Transformation

So categories include elements and morphisms. Functors consist of mappings from the elements and morphisms of one category to those of another (or the same) category. **Natural transformations** are a third level of mappings, from one functor to another.

Where `G`

and `H`

are functors from category **C** to category **D**, a natural transformation η between `G`

and `H`

is a family of morphisms `η[C1]:G(C1)→H(C1)`

in **D** for each element `C1`

of **C**. That is, `η[C1]`

has as source `C1`

's image under `G`

in **D**, and as target `C1`

's image under `H`

in **D**. The morphisms in this family must also satisfy the constraint:

for every morphism f:C1→C2 inC: η[C2] ∘ G(f) = H(f) ∘ η[C1]

That is, the morphism via `G(f)`

from `G(C1)`

to `G(C2)`

, and then via `η[C2]`

to `H(C2)`

, is identical to the morphism from `G(C1)`

via `η[C1]`

to `H(C1)`

, and then via `H(f)`

from `H(C1)`

to `H(C2)`

.

How natural transformations compose:

Consider four categories **B**, **C**, **D**, and **E**. Let `F`

be a functor from **B** to **C**; `G`

, `H`

, and `J`

be functors from **C** to **D**; and `K`

and `L`

be functors from **D** to **E**. Let η be a natural transformation from `G`

to `H`

; φ be a natural transformation from `H`

to `J`

; and ψ be a natural transformation from `K`

to `L`

. Pictorally:

-B-+ +---C--+ +----D-----+ +--E-- | | | | | | F: ------> G: ------> K: ------> | | | | | η | | | ψ | | | | v | | v | | H: ------> L: ------> | | | | | φ | | | | | | v | | | | J: ------> | | -----+ +--------+ +------------+ +-------

Then `(η F)`

is a natural transformation from the (composite) functor `GF`

to the composite functor `HF`

, such that where `B1`

is an element of category **B**, `(η F)[B1] = η[F(B1)]`

---that is, the morphism in **D** that `η`

assigns to the element `F(B1)`

of **C**.

And `(K η)`

is a natural transformation from the (composite) functor `KG`

to the (composite) functor `KH`

, such that where `C1`

is an element of category **C**, `(K η)[C1] = K(η[C1])`

---that is, the morphism in **E** that `K`

assigns to the morphism `η[C1]`

of **D**.

`(φ -v- η)`

is a natural transformation from `G`

to `J`

; this is known as a "vertical composition". For any morphism `f:C1→C2`

in **C**:

φ[C2] ∘ H(f) ∘ η[C1] = φ[C2] ∘ H(f) ∘ η[C1]

by naturalness of `φ`

, is:

φ[C2] ∘ H(f) ∘ η[C1] = J(f) ∘ φ[C1] ∘ η[C1]

by naturalness of `η`

, is:

φ[C2] ∘ η[C2] ∘ G(f) = J(f) ∘ φ[C1] ∘ η[C1]

Hence, we can define `(φ -v- η)[_]`

as: `φ[_] ∘ η[_]`

and rely on it to satisfy the constraints for a natural transformation from `G`

to `J`

:

(φ -v- η)[C2] ∘ G(f) = J(f) ∘ (φ -v- η)[C1]

An observation we'll rely on later: given the definitions of vertical composition and of how natural transformations compose with functors, it follows that:

((φ -v- η) F) = ((φ F) -v- (η F))

I'll assert without proving that vertical composition is associative and has an identity, which we'll call "the identity transformation."

`(ψ -h- η)`

is natural transformation from the (composite) functor `KG`

to the (composite) functor `LH`

; this is known as a "horizontal composition." It's trickier to define, but we won't be using it here. For reference:

(φ -h- η)[C1] = L(η[C1]) ∘ ψ[G(C1)] = ψ[H(C1)] ∘ K(η[C1])

Horizontal composition is also associative, and has the same identity as vertical composition.

## Monads

In earlier days, these were also called "triples."

A **monad** is a structure consisting of an (endo)functor `M`

from some category **C** to itself, along with some natural transformations, which we'll specify in a moment.

Let `T`

be a set of natural transformations `φ`

, each being between some arbitrary endofunctor `F`

on **C** and another functor which is the composite `MF'`

of `M`

and another arbitrary endofunctor `F'`

on **C**. That is, for each element `C1`

in **C**, `φ`

assigns `C1`

a morphism from element `F(C1)`

to element `MF'(C1)`

, satisfying the constraints detailed in the previous section. For different members of `T`

, the relevant functors may differ; that is, `φ`

is a transformation from functor `F`

to `MF'`

, `γ`

is a transformation from functor `G`

to `MG'`

, and none of `F`

, `F'`

, `G`

, `G'`

need be the same.

One of the members of `T`

will be designated the `unit`

transformation for `M`

, and it will be a transformation from the identity functor `1C`

for **C** to `M(1C)`

. So it will assign to `C1`

a morphism from `C1`

to `M(C1)`

.

We also need to designate for `M`

a `join`

transformation, which is a natural transformation from the (composite) functor `MM`

to `M`

.

These two natural transformations have to satisfy some constraints ("the monad laws") which are most easily stated if we can introduce a defined notion.

Let `φ`

and `γ`

be members of `T`

, that is they are natural transformations from `F`

to `MF'`

and from `G`

to `MG'`

, respectively. Let them be such that `F' = G`

. Now `(M γ)`

will also be a natural transformation, formed by composing the functor `M`

with the natural transformation `γ`

. Similarly, `(join G')`

will be a natural transformation, formed by composing the natural transformation `join`

with the functor `G'`

; it will transform the functor `MMG'`

to the functor `MG'`

. Now take the vertical composition of the three natural transformations `(join G')`

, `(M γ)`

, and `φ`

, and abbreviate it as follows. Since composition is associative I don't specify the order of composition on the rhs.

γ <=< φ =def. ((join G') -v- (M γ) -v- φ)

In other words, `<=<`

is a binary operator that takes us from two members `φ`

and `γ`

of `T`

to a composite natural transformation. (In functional programming, at least, this is called the "Kleisli composition operator". Sometimes it's written `φ >=> γ`

where that's the same as `γ <=< φ`

.)

`φ`

is a transformation from `F`

to `MF'`

, where the latter = `MG`

; `(M γ)`

is a transformation from `MG`

to `MMG'`

; and `(join G')`

is a transformation from `MMG'`

to `MG'`

. So the composite `γ <=< φ`

will be a transformation from `F`

to `MG'`

, and so also eligible to be a member of `T`

.

Now we can specify the "monad laws" governing a monad as follows:

(T, <=<, unit) constitute a monoid

That's it. Well, there may be a wrinkle here. I don't know whether the definition of a monoid requires the operation to be defined for every pair in its set. In the present case, `γ <=< φ`

isn't fully defined on `T`

, but only when `φ`

is a transformation to some `MF'`

and `γ`

is a transformation from `F'`

. But wherever `<=<`

is defined, the monoid laws must hold:

(i) γ <=< φ is also in T (ii) (ρ <=< γ) <=< φ = ρ <=< (γ <=< φ) (iii.1) unit <=< φ = φ (here φ has to be a natural transformation to M(1C)) (iii.2) ρ = ρ <=< unit (here ρ has to be a natural transformation from 1C)

If `φ`

is a natural transformation from `F`

to `M(1C)`

and `γ`

is `(φ G')`

, that is, a natural transformation from `FG'`

to `MG'`

, then we can extend (iii.1) as follows:

γ = (φ G') = ((unit <=< φ) G') since unit is a natural transformation to M(1C), this is: = (((join 1C) -v- (M unit) -v- φ) G') = (((join 1C) G') -v- ((M unit) G') -v- (φ G')) = ((join (1C G')) -v- (M (unit G')) -v- γ) = ((join G') -v- (M (unit G')) -v- γ) since (unit G') is a natural transformation to MG', this is: = (unit G') <=< γ

where as we said `γ`

is a natural transformation from some `FG'`

to `MG'`

.

Similarly, if `ρ`

is a natural transformation from `1C`

to `MR'`

, and `γ`

is `(ρ G)`

, that is, a natural transformation from `G`

to `MR'G`

, then we can extend (iii.2) as follows:

γ = (ρ G) = ((ρ <=< unit) G) = since ρ is a natural transformation to MR', this is: = (((join R') -v- (M ρ) -v- unit) G) = (((join R') G) -v- ((M ρ) G) -v- (unit G)) = ((join (R'G)) -v- (M (ρ G)) -v- (unit G)) since γ = (ρ G) is a natural transformation to MR'G, this is: = γ <=< (unit G)

where as we said `γ`

is a natural transformation from `G`

to some `MR'G`

.

Summarizing then, the monad laws can be expressed as:

For all ρ, γ, φ in T for which ρ <=< γ and γ <=< φ are defined: (i) γ <=< φ etc are also in T (ii) (ρ <=< γ) <=< φ = ρ <=< (γ <=< φ) (iii.1) (unit G') <=< γ = γ whenever γ is a natural transformation from some FG' to MG' (iii.2) γ = γ <=< (unit G) whenever γ is a natural transformation from G to some MR'G

## Getting to the standard category-theory presentation of the monad laws

In category theory, the monad laws are usually stated in terms of `unit`

and `join`

instead of `unit`

and `<=<`

.

Let's remind ourselves of principles stated above:

composition of morphisms, functors, and natural compositions is associative

functors "distribute over composition", that is for any morphisms

`f`

and`g`

in`F`

's source category:`F(g ∘ f) = F(g) ∘ F(f)`

if

`η`

is a natural transformation from`G`

to`H`

, then for every`f:C1→C2`

in`G`

and`H`

's source category**C**:`η[C2] ∘ G(f) = H(f) ∘ η[C1]`

.`(η F)[X] = η[F(X)]`

`(K η)[X] = K(η[X])`

`((φ -v- η) F) = ((φ F) -v- (η F))`

Let's use the definitions of naturalness, and of composition of natural transformations, to establish two lemmas.

Recall that `join`

is a natural transformation from the (composite) functor `MM`

to `M`

. So for elements `C1`

in **C**, `join[C1]`

will be a morphism from `MM(C1)`

to `M(C1)`

. And for any morphism `f:C1→C2`

in **C**:

(1) join[C2] ∘ MM(f) = M(f) ∘ join[C1]

Next, let `γ`

be a transformation from `G`

to `MG'`

, and
consider the composite transformation `((join MG') -v- (MM γ))`

.

`γ`

assigns elements`C1`

in**C**a morphism`γ*:G(C1) → MG'(C1)`

.`(MM γ)`

is a transformation that instead assigns`C1`

the morphism`MM(γ*)`

.`(join MG')`

is a transformation from`MM(MG')`

to`M(MG')`

that assigns`C1`

the morphism`join[MG'(C1)]`

.

Composing them:

(2) ((join MG') -v- (MM γ)) assigns to C1 the morphism join[MG'(C1)] ∘ MM(γ*).

Next, consider the composite transformation `((M γ) -v- (join G))`

:

(3) ((M γ) -v- (join G)) assigns to C1 the morphism M(γ*) ∘ join[G(C1)].

So for every element `C1`

of **C**:

((join MG') -v- (MM γ))[C1], by (2) is: join[MG'(C1)] ∘ MM(γ*), which by (1), with f=γ*:G(C1)→MG'(C1) is: M(γ*) ∘ join[G(C1)], which by 3 is: ((M γ) -v- (join G))[C1]

So our **(lemma 1)** is:

((join MG') -v- (MM γ)) = ((M γ) -v- (join G)), where as we said γ is a natural transformation from G to MG'.

Next recall that `unit`

is a natural transformation from `1C`

to `M`

. So for elements `C1`

in **C**, `unit[C1]`

will be a morphism from `C1`

to `M(C1)`

. And for any morphism `f:C1→C2`

in **C**:

(4) unit[C2] ∘ f = M(f) ∘ unit[C1]

Next, consider the composite transformation `((M γ) -v- (unit G))`

:

(5) ((M γ) -v- (unit G)) assigns to C1 the morphism M(γ*) ∘ unit[G(C1)].

Next, consider the composite transformation `((unit MG') -v- γ)`

:

(6) ((unit MG') -v- γ) assigns to C1 the morphism unit[MG'(C1)] ∘ γ*.

So for every element C1 of **C**:

((M γ) -v- (unit G))[C1], by (5) = M(γ*) ∘ unit[G(C1)], which by (4), with f=γ*:G(C1)→MG'(C1) is: unit[MG'(C1)] ∘ γ*, which by (6) = ((unit MG') -v- γ)[C1]

So our **(lemma 2)** is:

(((M γ) -v- (unit G)) = ((unit MG') -v- γ)), where as we said γ is a natural transformation from G to MG'.

Finally, we substitute `((join G') -v- (M γ) -v- φ)`

for `γ <=< φ`

in the monad laws. For simplicity, I'll omit the "-v-".

For all ρ, γ, φ in T, where φ is a transformation from F to MF', γ is a transformation from G to MG', ρ is a transformation from R to MR', and F'=G and G'=R: (i) γ <=< φ etc are also in T ==> (i') ((join G') (M γ) φ) etc are also in T

(ii) (ρ <=< γ) <=< φ = ρ <=< (γ <=< φ) ==> (ρ <=< γ) is a transformation from G to MR', so (ρ <=< γ) <=< φ becomes: ((join R') (M (ρ <=< γ)) φ) which is: ((join R') (M ((join R') (M ρ) γ)) φ) similarly, ρ <=< (γ <=< φ) is: ((join R') (M ρ) ((join G') (M γ) φ)) substituting these into (ii), and helping ourselves to associativity on the rhs, we get: ((join R') (M ((join R') (M ρ) γ)) φ) = ((join R') (M ρ) (join G') (M γ) φ) which by the distributivity of functors over composition, and helping ourselves to associativity on the lhs, yields: ((join R') (M join R') (MM ρ) (M γ) φ) = ((join R') (M ρ) (join G') (M γ) φ) which by lemma 1, with ρ a transformation from G' to MR', yields: ((join R') (M join R') (MM ρ) (M γ) φ) = ((join R') (join MR') (MM ρ) (M γ) φ) [-- Are the next two steps too cavalier? --] which will be true for all ρ, γ, φ only when: ((join R') (M join R')) = ((join R') (join MR')), for any R' which will in turn be true when: (ii') (join (M join)) = (join (join M))

(iii.1) (unit G') <=< γ = γ when γ is a natural transformation from some FG' to MG' ==> (unit G') is a transformation from G' to MG', so: (unit G') <=< γ becomes: ((join G') (M (unit G')) γ) which is: ((join G') ((M unit) G') γ) substituting in (iii.1), we get: ((join G') ((M unit) G') γ) = γ which is: (((join (M unit)) G') γ) = γ [-- Are the next two steps too cavalier? --] which will be true for all γ just in case: for any G', ((join (M unit)) G') = the identity transformation which will in turn be true just in case: (iii.1') (join (M unit)) = the identity transformation

(iii.2) γ = γ <=< (unit G) when γ is a natural transformation from G to some MR'G ==> γ <=< (unit G) becomes: ((join R'G) (M γ) (unit G)) substituting in (iii.2), we get: γ = ((join R'G) (M γ) (unit G)) which by lemma 2, yields: γ = (((join R'G) ((unit MR'G) γ) which is: γ = (((join (unit M)) R'G) γ) [-- Are the next two steps too cavalier? --] which will be true for all γ just in case: for any R'G, ((join (unit M)) R'G) = the identity transformation which will in turn be true just in case: (iii.2') (join (unit M)) = the identity transformation

Collecting the results, our monad laws turn out in this format to be:

For all ρ, γ, φ in T, where φ is a transformation from F to MF', γ is a transformation from G to MG', ρ is a transformation from R to MR', and F'=G and G'=R: (i') ((join G') (M γ) φ) etc also in T (ii') (join (M join)) = (join (join M)) (iii.1') (join (M unit)) = the identity transformation (iii.2') (join (unit M)) = the identity transformation

In category-theory presentations, you may see `unit`

referred to as `η`

, and `join`

referred to as `μ`

. Also, instead of the monad `(M, unit, join)`

, you may sometimes see discussion of the "Kleisli triple" `(M, unit, =<<)`

. Alternatively, `=<<`

may be called `⋆`

. These are interdefinable (see below).

## Getting to the functional programming presentation of the monad laws

In functional programming, `unit`

is sometimes called `return`

and the monad laws are usually stated in terms of `unit`

/`return`

and an operation called `bind`

which is interdefinable with `<=<`

or with `join`

.

The base category **C** will have types as elements, and monadic functions as its morphisms. The source and target of a morphism will be the types of its argument and its result. (As always, there can be multiple distinct morphisms from the same source to the same target.)

A monad `M`

will consist of a mapping from types `'t`

to types `M('t)`

, and a mapping from functions `f:C1→C2`

to functions `M(f):M(C1)→M(C2)`

. This is also known as `lift`

for _{M} f`M`

, and is pronounced "function f lifted into the monad M." For example, where `M`

is the List monad, `M`

maps every type `'t`

into the type `'t list`

, and maps every function `f:x→y`

into the function that maps `[x1,x2...]`

to `[y1,y2,...]`

.

In functional programming, instead of working with natural transformations we work with "monadic values" and polymorphic functions "into the monad."

A "monadic value" is any member of a type `M('t)`

, for any type `'t`

. For example, any `int list`

is a monadic value for the List monad. We can think of these monadic values as the result of applying some function `phi`

, whose type is `F('t) -> M(F'('t))`

. `'t`

here is any collection of free type variables, and `F('t)`

and `F'('t)`

are types parameterized on `'t`

. An example, with `M`

being the List monad, `'t`

being `('t1,'t2)`

, `F('t1,'t2)`

being `char * 't1 * 't2`

, and `F'('t1,'t2)`

being `int * 't1 * 't2`

:

let phi = fun ((_:char), x, y) -> [(1,x,y),(2,x,y)]

[-- I intentionally chose this polymorphic function because simpler ways of mapping the polymorphic monad operations from functional programming onto the category theory notions can't accommodate it. We have all the F, MF' (unit G') and so on in order to be able to be handle even phis like this. --]

Now where `gamma`

is another function of type `F'('t) -> M(G'('t))`

, we define:

gamma =<< phi a =def. ((join G') -v- (M gamma)) (phi a) = ((join G') -v- (M gamma) -v- phi) a = (gamma <=< phi) a

Hence:

gamma <=< phi = (fun a -> gamma =<< phi a)

`gamma =<< phi a`

is called the operation of "binding" the function gamma to the monadic value `phi a`

, and is usually written as `phi a >>= gamma`

.

With these definitions, our monadic laws become:

Where phi is a polymorphic function of type F('t) -> M(F'('t)) gamma is a polymorphic function of type G('t) -> M(G'('t)) rho is a polymorphic function of type R('t) -> M(R'('t)) and F' = G and G' = R, and a ranges over values of type F('t), and b ranges over values of type G('t), and c ranges over values of type G'('t): (i) γ <=< φ is defined, and is a natural transformation from F to MG' ==> (i'') fun a -> gamma =<< phi a is defined, and is a function from type F('t) -> M(G'('t))

(ii) (ρ <=< γ) <=< φ = ρ <=< (γ <=< φ) ==> (fun a -> (rho <=< gamma) =<< phi a) = (fun a -> rho =<< (gamma <=< phi) a) (fun a -> (fun b -> rho =<< gamma b) =<< phi a) = (fun a -> rho =<< (gamma =<< phi a)) (ii'') (fun b -> rho =<< gamma b) =<< phi a = rho =<< (gamma =<< phi a)

(iii.1) (unit G') <=< γ = γ whenever γ is a natural transformation from some FG' to MG' ==> (unit G') <=< gamma = gamma whenever gamma is a function of type F(G'('t)) -> M(G'('t)) (fun b -> (unit G') =<< gamma b) = gamma (unit G') =<< gamma b = gamma b Let return be a polymorphic function mapping arguments of any type 't to M('t). In particular, it maps arguments c of type G'('t) to the monadic value (unit G') c, of type M(G'('t)). (iii.1'') return =<< gamma b = gamma b

(iii.2) γ = γ <=< (unit G) whenever γ is a natural transformation from G to some MR'G ==> gamma = gamma <=< (unit G) whenever gamma is a function of type G('t) -> M(R'(G('t))) gamma = (fun b -> gamma =<< (unit G) b) As above, return will map arguments b of type G('t) to the monadic value (unit G) b, of type M(G('t)). gamma = (fun b -> gamma =<< return b) (iii.2'') gamma b = gamma =<< return b

Summarizing (ii''), (iii.1''), (iii.2''), these are the monadic laws as usually stated in the functional programming literature:

`(fun b -> rho =<< gamma b) =<< phi a = rho =<< (gamma =<< phi a)`

Usually written reversed, and with a monadic variable

`u`

standing in for`phi a`

:`u >>= (fun b -> gamma b >>= rho) = (u >>= gamma) >>= rho`

`return =<< gamma b = gamma b`

Usually written reversed, and with

`u`

standing in for`gamma b`

:`u >>= return = u`

`gamma b = gamma =<< return b`

Usually written reversed:

`return b >>= gamma = gamma b`