RecordsAdding Records to STLC

Set Warnings "-notation-overridden,-parsing".
From Coq Require Import Strings.String.
From PLF Require Import Maps.
From PLF Require Import Imp.
From PLF Require Import Smallstep.
From PLF Require Import Stlc.

Adding Records

We saw in chapter MoreStlc how records can be treated as just syntactic sugar for nested uses of products. This is OK for simple examples, but the encoding is informal (in reality, if we actually treated records this way, it would be carried out in the parser, which we are eliding here), and anyway it is not very efficient. So it is also interesting to see how records can be treated as first-class citizens of the language. This chapter shows how.
Recall the informal definitions we gave before:
       t ::=                          Terms:
           | {i1=t1, ..., in=tn}         record
           | t.i                         projection
           | ...

       v ::=                          Values:
           | {i1=v1, ..., in=vn}         record value
           | ...

       T ::=                          Types:
           | {i1:T1, ..., in:Tn}         record type
           | ...
ti ==> ti' (ST_Rcd)  

{i1=v1, ..., im=vm, in=tn, ...} ==> {i1=v1, ..., im=vm, in=tn', ...}
t1 ==> t1' (ST_Proj1)  

t1.i ==> t1'.i

{..., i=vi, ...}.i ==> vi
Gamma ⊢ t1 : T1     ...     Gamma ⊢ tn : Tn (T_Rcd)  

Gamma ⊢ {i1=t1, ..., in=tn} : {i1:T1, ..., in:Tn}
Gamma ⊢ t : {..., i:Ti, ...} (T_Proj)  

Gamma ⊢ t.i : Ti

Formalizing Records

Module STLCExtendedRecords.

Syntax and Operational Semantics

The most obvious way to formalize the syntax of record types would be this:
Module FirstTry.

Definition alist (X : Type) := list (string * X).

Inductive ty : Type :=
  | Base : stringty
  | Arrow : tytyty
  | TRcd : (alist ty) → ty.
Unfortunately, we encounter here a limitation in Coq: this type does not automatically give us the induction principle we expect: the induction hypothesis in the TRcd case doesn't give us any information about the ty elements of the list, making it useless for the proofs we want to do.
(* Check ty_ind.
    ty_ind :
      forall P : ty -> Prop,
        (forall i : id, P (Base i)) ->
        (forall t : ty, P t -> forall t0 : ty, P t0 
                            -> P (Arrow t t0)) ->
        (forall a : alist ty, P (TRcd a)) ->    (* ??? *)
        forall t : ty, P t

End FirstTry.
It is possible to get a better induction principle out of Coq, but the details of how this is done are not very pretty, and the principle we obtain is not as intuitive to use as the ones Coq generates automatically for simple Inductive definitions.
Fortunately, there is a different way of formalizing records that is, in some ways, even simpler and more natural: instead of using the standard Coq list type, we can essentially incorporate its constructors ("nil" and "cons") in the syntax of our types.
Inductive ty : Type :=
  | Base : stringty
  | Arrow : tytyty
  | RNil : ty
  | RCons : stringtytyty.
Similarly, at the level of terms, we have constructors trnil, for the empty record, and rcons, which adds a single field to the front of a list of fields.
Inductive tm : Type :=
  | var : stringtm
  | app : tmtmtm
  | abs : stringtytmtm
  (* records *)
  | rproj : tmstringtm
  | trnil : tm
  | rcons : stringtmtmtm.
Some examples...
Open Scope string_scope.

Notation a := "a".
Notation f := "f".
Notation g := "g".
Notation l := "l".
Notation A := (Base "A").
Notation B := (Base "B").
Notation k := "k".
Notation i1 := "i1".
Notation i2 := "i2".
{ i1:A }
(* Check (RCons i1 A RNil). *)
{ i1:AB, i2:A }
(* Check (RCons i1 (Arrow A B)
           (RCons i2 A RNil)). *)


One issue with generalizing the abstract syntax for records from lists to the nil/cons presentation is that it introduces the possibility of writing strange types like this...
Definition weird_type := RCons X A B.
where the "tail" of a record type is not actually a record type!
We'll structure our typing judgement so that no ill-formed types like weird_type are ever assigned to terms. To support this, we define predicates record_ty and record_tm, which identify record types and terms, and well_formed_ty which rules out the ill-formed types.
First, a type is a record type if it is built with just RNil and RCons at the outermost level.
Inductive record_ty : tyProp :=
  | RTnil :
        record_ty RNil
  | RTcons : i T1 T2,
        record_ty (RCons i T1 T2).
With this, we can define well-formed types.
Inductive well_formed_ty : tyProp :=
  | wfBase : i,
        well_formed_ty (Base i)
  | wfArrow : T1 T2,
        well_formed_ty T1
        well_formed_ty T2
        well_formed_ty (Arrow T1 T2)
  | wfRNil :
        well_formed_ty RNil
  | wfRCons : i T1 T2,
        well_formed_ty T1
        well_formed_ty T2
        record_ty T2
        well_formed_ty (RCons i T1 T2).

Hint Constructors record_ty well_formed_ty.
Note that record_ty is not recursive — it just checks the outermost constructor. The well_formed_ty property, on the other hand, verifies that the whole type is well formed in the sense that the tail of every record (the second argument to RCons) is a record.
Of course, we should also be concerned about ill-formed terms, not just types; but typechecking can rule those out without the help of an extra well_formed_tm definition because it already examines the structure of terms. All we need is an analog of record_ty saying that a term is a record term if it is built with trnil and rcons.
Inductive record_tm : tmProp :=
  | rtnil :
        record_tm trnil
  | rtcons : i t1 t2,
        record_tm (rcons i t1 t2).

Hint Constructors record_tm.


Substitution extends easily.
Fixpoint subst (x:string) (s:tm) (t:tm) : tm :=
  match t with
  | var yif eqb_string x y then s else t
  | abs y T t1abs y T
                     (if eqb_string x y then t1 else (subst x s t1))
  | app t1 t2app (subst x s t1) (subst x s t2)
  | rproj t1 irproj (subst x s t1) i
  | trniltrnil
  | rcons i t1 tr1rcons i (subst x s t1) (subst x s tr1)

Notation "'[' x ':=' s ']' t" := (subst x s t) (at level 20).


A record is a value if all of its fields are.
Inductive value : tmProp :=
  | v_abs : x T11 t12,
      value (abs x T11 t12)
  | v_rnil : value trnil
  | v_rcons : i v1 vr,
      value v1
      value vr
      value (rcons i v1 vr).

Hint Constructors value.
To define reduction, we'll need a utility function for extracting one field from record term:
Fixpoint tlookup (i:string) (tr:tm) : option tm :=
  match tr with
  | rcons i' t tr'if eqb_string i i' then Some t else tlookup i tr'
  | _None
The step function uses this term-level lookup function in the projection rule.
Reserved Notation "t1 '-->' t2" (at level 40).

Inductive step : tmtmProp :=
  | ST_AppAbs : x T11 t12 v2,
         value v2
         (app (abs x T11 t12) v2) --> ([x:=v2]t12)
  | ST_App1 : t1 t1' t2,
         t1 --> t1'
         (app t1 t2) --> (app t1' t2)
  | ST_App2 : v1 t2 t2',
         value v1
         t2 --> t2'
         (app v1 t2) --> (app v1 t2')
  | ST_Proj1 : t1 t1' i,
        t1 --> t1'
        (rproj t1 i) --> (rproj t1' i)
  | ST_ProjRcd : tr i vi,
        value tr
        tlookup i tr = Some vi
        (rproj tr i) --> vi
  | ST_Rcd_Head : i t1 t1' tr2,
        t1 --> t1'
        (rcons i t1 tr2) --> (rcons i t1' tr2)
  | ST_Rcd_Tail : i v1 tr2 tr2',
        value v1
        tr2 --> tr2'
        (rcons i v1 tr2) --> (rcons i v1 tr2')

where "t1 '-->' t2" := (step t1 t2).

Notation multistep := (multi step).
Notation "t1 '-->*' t2" := (multistep t1 t2) (at level 40).

Hint Constructors step.


Next we define the typing rules. These are nearly direct transcriptions of the inference rules shown above: the only significant difference is the use of well_formed_ty. In the informal presentation we used a grammar that only allowed well-formed record types, so we didn't have to add a separate check.
One sanity condition that we'd like to maintain is that, whenever has_type Gamma t T holds, will also be the case that well_formed_ty T, so that has_type never assigns ill-formed types to terms. In fact, we prove this theorem below. However, we don't want to clutter the definition of has_type with unnecessary uses of well_formed_ty. Instead, we place well_formed_ty checks only where needed: where an inductive call to has_type won't already be checking the well-formedness of a type. For example, we check well_formed_ty T in the T_Var case, because there is no inductive has_type call that would enforce this. Similarly, in the T_Abs case, we require a proof of well_formed_ty T11 because the inductive call to has_type only guarantees that T12 is well-formed.
Fixpoint Tlookup (i:string) (Tr:ty) : option ty :=
  match Tr with
  | RCons i' T Tr'
      if eqb_string i i' then Some T else Tlookup i Tr'
  | _None

Definition context := partial_map ty.

Reserved Notation "Gamma '⊢' t '∈' T" (at level 40).

Inductive has_type : contexttmtyProp :=
  | T_Var : Gamma x T,
      Gamma x = Some T
      well_formed_ty T
      Gamma ⊢ (var x) ∈ T
  | T_Abs : Gamma x T11 T12 t12,
      well_formed_ty T11
      (update Gamma x T11) ⊢ t12T12
      Gamma ⊢ (abs x T11 t12) ∈ (Arrow T11 T12)
  | T_App : T1 T2 Gamma t1 t2,
      Gammat1 ∈ (Arrow T1 T2) →
      Gamma ⊢ (app t1 t2) ∈ T2
  (* records: *)
  | T_Proj : Gamma i t Ti Tr,
      Tlookup i Tr = Some Ti
      Gamma ⊢ (rproj t i) ∈ Ti
  | T_RNil : Gamma,
  | T_RCons : Gamma i t T tr Tr,
      record_ty Tr
      record_tm tr
      Gamma ⊢ (rcons i t tr) ∈ (RCons i T Tr)

where "Gamma '⊢' t '∈' T" := (has_type Gamma t T).

Hint Constructors has_type.


Exercise: 2 stars, standard (examples)

Finish the proofs below. Feel free to use Coq's automation features in this proof. However, if you are not confident about how the type system works, you may want to carry out the proofs first using the basic features (apply instead of eapply, in particular) and then perhaps compress it using automation. Before starting to prove anything, make sure you understand what it is saying.
Lemma typing_example_2 :
    (app (abs a (RCons i1 (Arrow A A)
                      (RCons i2 (Arrow B B)
              (rproj (var a) i2))
            (rcons i1 (abs a A (var a))
            (rcons i2 (abs a B (var a))
             trnil))) ∈
    (Arrow B B).
  (* FILL IN HERE *) Admitted.

Example typing_nonexample :
      (update empty a (RCons i2 (Arrow A A)
                                RNil)) ⊢
               (rcons i1 (abs a B (var a)) (var a)) ∈
  (* FILL IN HERE *) Admitted.

Example typing_nonexample_2 : y,
    (update empty y A) ⊢
           (app (abs a (RCons i1 A RNil)
                     (rproj (var a) i1))
                   (rcons i1 (var y) (rcons i2 (var y) trnil))) ∈
  (* FILL IN HERE *) Admitted.

Properties of Typing

The proofs of progress and preservation for this system are essentially the same as for the pure simply typed lambda-calculus, but we need to add some technical lemmas involving records.


Lemma wf_rcd_lookup : i T Ti,
  well_formed_ty T
  Tlookup i T = Some Ti
  well_formed_ty Ti.
Proof with eauto.
  intros i T.
  induction T; intros; try solve_by_invert.
  - (* RCons *)
    inversion H. subst. unfold Tlookup in H0.
    destruct (eqb_string i s)...
    inversion H0. subst... Qed.

Lemma step_preserves_record_tm : tr tr',
  record_tm tr
  tr --> tr'
  record_tm tr'.
  intros tr tr' Hrt Hstp.
  inversion Hrt; subst; inversion Hstp; subst; auto.

Lemma has_type__wf : Gamma t T,
  GammatTwell_formed_ty T.
Proof with eauto.
  intros Gamma t T Htyp.
  induction Htyp...
  - (* T_App *)
    inversion IHHtyp1...
  - (* T_Proj *)
    eapply wf_rcd_lookup...

Field Lookup

Lemma: If empty v : T and Tlookup i T returns Some Ti, then tlookup i v returns Some ti for some term ti such that empty ti Ti.
Proof: By induction on the typing derivation Htyp. Since Tlookup i T = Some Ti, T must be a record type, this and the fact that v is a value eliminate most cases by inspection, leaving only the T_RCons case.
If the last step in the typing derivation is by T_RCons, then t = rcons i0 t tr and T = RCons i0 T Tr for some i0, t, tr, T and Tr.
This leaves two possiblities to consider - either i0 = i or not.
  • If i = i0, then since Tlookup i (RCons i0 T Tr) = Some Ti we have T = Ti. It follows that t itself satisfies the theorem.
  • On the other hand, suppose i i0. Then
    Tlookup i T = Tlookup i Tr
    tlookup i t = tlookup i tr,
    so the result follows from the induction hypothesis.
Here is the formal statement:
Lemma lookup_field_in_value : v T i Ti,
  value v
  Tlookup i T = Some Ti
  ti, tlookup i v = Some tiemptytiTi.
Proof with eauto.
  intros v T i Ti Hval Htyp Hget.
  remember (@empty ty) as Gamma.
  induction Htyp; subst; try solve_by_invert...
  - (* T_RCons *)
    simpl in Hget. simpl. destruct (eqb_string i i0).
    + (* i is first *)
      simpl. inversion Hget. subst.
    + (* get tail *)
      destruct IHHtyp2 as [vi [Hgeti Htypi]]...
      inversion Hval... Qed.


Theorem progress : t T,
     value tt', t --> t'.
Proof with eauto.
  (* Theorem: Suppose empty ⊢ t : T.  Then either
       1. t is a value, or
       2. t --> t' for some t'.
     Proof: By induction on the given typing derivation. *)

  intros t T Ht.
  remember (@empty ty) as Gamma.
  generalize dependent HeqGamma.
  induction Ht; intros HeqGamma; subst.
  - (* T_Var *)
    (* The final rule in the given typing derivation cannot be 
       T_Var, since it can never be the case that 
       empty x : T (since the context is empty). *)

    inversion H.
  - (* T_Abs *)
    (* If the T_Abs rule was the last used, then 
       t = abs x T11 t12, which is a value. *)

  - (* T_App *)
    (* If the last rule applied was T_App, then t = t1 t2
       and we know from the form of the rule that
         empty t1 : T1 T2
         empty t2 : T1
       By the induction hypothesis, each of t1 and t2 either is a value
       or can take a step. *)

    destruct IHHt1; subst...
    + (* t1 is a value *)
      destruct IHHt2; subst...
      * (* t2 is a value *)
      (* If both t1 and t2 are values, then we know that
         t1 = abs x T11 t12, since abstractions are the only 
         values that can have an arrow type.  But
         (abs x T11 t12) t2 --> [x:=t2]t12 by ST_AppAbs. *)

        inversion H; subst; try solve_by_invert.
      * (* t2 steps *)
        (* If t1 is a value and t2 --> t2', then
           t1 t2 --> t1 t2' by ST_App2. *)

        destruct H0 as [t2' Hstp]. (app t1 t2')...
    + (* t1 steps *)
      (* Finally, If t1 --> t1', then t1 t2 --> t1' t2
         by ST_App1. *)

      destruct H as [t1' Hstp]. (app t1' t2)...
  - (* T_Proj *)
    (* If the last rule in the given derivation is T_Proj, then
       t = rproj t i and
           empty t : (TRcd Tr)
       By the IH, t either is a value or takes a step. *)

    right. destruct IHHt...
    + (* rcd is value *)
      (* If t is a value, then we may use lemma
         lookup_field_in_value to show tlookup i t = Some ti 
         for some ti which gives us rproj i t --> ti by
         ST_ProjRcd. *)

      destruct (lookup_field_in_value _ _ _ _ H0 Ht H)
        as [ti [Hlkup _]].
    + (* rcd_steps *)
      (* On the other hand, if t --> t', then
         rproj t i --> rproj t' i by ST_Proj1. *)

      destruct H0 as [t' Hstp]. (rproj t' i)...
  - (* T_RNil *)
    (* If the last rule in the given derivation is T_RNil
       then t = trnil, which is a value. *)

  - (* T_RCons *)
    (* If the last rule is T_RCons, then t = rcons i t tr and
         empty t : T
         empty tr : Tr
       By the IH, each of t and tr either is a value or can 
       take a step. *)

    destruct IHHt1...
    + (* head is a value *)
      destruct IHHt2; try reflexivity.
      * (* tail is a value *)
      (* If t and tr are both values, then rcons i t tr
         is a value as well. *)

      * (* tail steps *)
        (* If t is a value and tr --> tr', then
           rcons i t tr --> rcons i t tr' by
           ST_Rcd_Tail. *)

        right. destruct H2 as [tr' Hstp].
        (rcons i t tr')...
    + (* head steps *)
      (* If t --> t', then
         rcons i t tr --> rcons i t' tr
         by ST_Rcd_Head. *)

      right. destruct H1 as [t' Hstp].
      (rcons i t' tr)... Qed.

Context Invariance

Inductive appears_free_in : stringtmProp :=
  | afi_var : x,
      appears_free_in x (var x)
  | afi_app1 : x t1 t2,
      appears_free_in x t1appears_free_in x (app t1 t2)
  | afi_app2 : x t1 t2,
      appears_free_in x t2appears_free_in x (app t1 t2)
  | afi_abs : x y T11 t12,
        appears_free_in x t12
        appears_free_in x (abs y T11 t12)
  | afi_proj : x t i,
     appears_free_in x t
     appears_free_in x (rproj t i)
  | afi_rhead : x i ti tr,
      appears_free_in x ti
      appears_free_in x (rcons i ti tr)
  | afi_rtail : x i ti tr,
      appears_free_in x tr
      appears_free_in x (rcons i ti tr).

Hint Constructors appears_free_in.

Lemma context_invariance : Gamma Gamma' t S,
     (x, appears_free_in x tGamma x = Gamma' x) →
Proof with eauto.
  intros. generalize dependent Gamma'.
  induction H;
    intros Gamma' Heqv...
  - (* T_Var *)
    apply T_Var... rewrite <- Heqv...
  - (* T_Abs *)
    apply T_Abs... apply IHhas_type. intros y Hafi.
    unfold update, t_update. destruct (eqb_stringP x y)...
  - (* T_App *)
    apply T_App with T1...
  - (* T_RCons *)
    apply T_RCons... Qed.

Lemma free_in_context : x t T Gamma,
   appears_free_in x t
   T', Gamma x = Some T'.
Proof with eauto.
  intros x t T Gamma Hafi Htyp.
  induction Htyp; inversion Hafi; subst...
  - (* T_Abs *)
    destruct IHHtyp as [T' Hctx]... T'.
    unfold update, t_update in Hctx.
    rewrite false_eqb_string in Hctx...


Lemma substitution_preserves_typing : Gamma x U v t S,
     (update Gamma x U) ⊢ tS
     Gamma ⊢ ([x:=v]t) ∈ S.
Proof with eauto.
  (* Theorem: If x>U;Gamma ⊢ t : S and empty ⊢ v : U, then
     Gamma ⊢ (x:=vt) S. *)

  intros Gamma x U v t S Htypt Htypv.
  generalize dependent Gamma. generalize dependent S.
  (* Proof: By induction on the term t.  Most cases follow 
     directly from the IH, with the exception of var, 
     abs, rcons. The former aren't automatic because we 
     must reason about how the variables interact. In the 
     case of rcons, we must do a little extra work to show 
     that substituting into a term doesn't change whether 
     it is a record term. *)

  induction t;
    intros S Gamma Htypt; simpl; inversion Htypt; subst...
  - (* var *)
    simpl. rename s into y.
    (* If t = y, we know that
         empty v : U and
         x>U; Gamma y : S
       and, by inversion, update Gamma x U y = Some S.  
       We want to show that Gamma [x:=v]y : S.

       There are two cases to consider: either x=y or xy. *)

    unfold update, t_update in H0.
    destruct (eqb_stringP x y) as [Hxy|Hxy].
    + (* x=y *)
    (* If x = y, then we know that U = S, and that 
       [x:=v]y = v. So what we really must show is that 
       if empty v : U then Gamma v : U.  We have
        already proven a more general version of this theorem, 
        called context invariance! *)

      inversion H0; subst. clear H0.
      eapply context_invariance...
      intros x Hcontra.
      destruct (free_in_context _ _ S empty Hcontra)
        as [T' HT']...
      inversion HT'.
    + (* x<>y *)
    (* If x y, then Gamma y = Some S and the substitution
       has no effect.  We can show that Gamma y : S by 
       T_Var. *)

      apply T_Var...
  - (* abs *)
    rename s into y. rename t into T11.
    (* If t = abs y T11 t0, then we know that
         x>U; Gamma abs y T11 t0 : T11T12
         x>U; y>T11; Gamma t0 : T12
         empty v : U
       As our IH, we know that forall S Gamma,
         x>U; Gamma t0 : S Gamma [x:=v]t0 S.

       We can calculate that
        [x:=v]t = abs y T11 (if eqb_string x y then t0 else [x:=v]t0) ,
       and we must show that Gamma [x:=v]t : T11T12.  We know
       we will do so using T_Abs, so it remains to be shown that:
         y>T11; Gamma if eqb_string x y then t0 else [x:=v]t0 : T12
       We consider two cases: x = y and x y. *)

    apply T_Abs...
    destruct (eqb_stringP x y) as [Hxy|Hxy].
    + (* x=y *)
      (* If x = y, then the substitution has no effect.  Context
         invariance shows that y:U,y:T11 and Gamma,y:T11 are
         equivalent.  Since t0 : T12 under the former context, 
         this is also the case under the latter. *)

      eapply context_invariance...
      intros x Hafi. unfold update, t_update.
      destruct (eqb_string y x)...
    + (* x<>y *)
      (* If x y, then the IH and context invariance allow 
         us to show that
           x>U; y>T11; Gamma t0 : T12       =>
           y>T11; x>U; Gamma t0 : T12       =>
           y>T11; Gamma [x:=v]t0 : T12 *)

      apply IHt. eapply context_invariance...
      intros z Hafi. unfold update, t_update.
      destruct (eqb_stringP y z)...
      subst. rewrite false_eqb_string...
  - (* rcons *)
    apply T_RCons... inversion H7; subst; simpl...

Theorem preservation : t t' T,
     t --> t'
Proof with eauto.
  intros t t' T HT.
  (* Theorem: If empty t : T and t --> t', then
     empty t' : T. *)

  remember (@empty ty) as Gamma. generalize dependent HeqGamma.
  generalize dependent t'.
  (* Proof: By induction on the given typing derivation.  
     Many cases are contradictory (T_VarT_Abs) or follow 
     directly from the IH (T_RCons).  We show just the 
     interesting ones. *)

  induction HT;
    intros t' HeqGamma HE; subst; inversion HE; subst...
  - (* T_App *)
    (* If the last rule used was T_App, then t = t1 t2
       and three rules could have been used to show t --> t':
       ST_App1ST_App2, and ST_AppAbs. In the first two 
       cases, the result follows directly from the IH. *)

    inversion HE; subst...
    + (* ST_AppAbs *)
      (* For the third case, suppose
           t1 = abs x T11 t12
           t2 = v2.  We must show that empty [x:=v2]t12 : T2.
         We know by assumption that
             empty abs x T11 t12 : T1T2
         and by inversion
             x:T1 t12 : T2
         We have already proven that substitution_preserves_typing and
             empty v2 : T1
         by assumption, so we are done. *)

      apply substitution_preserves_typing with T1...
      inversion HT1...
  - (* T_Proj *)
    (* If the last rule was T_Proj, then t = rproj t1 i.  
       Two rules could have caused t --> t'T_Proj1 and
       T_ProjRcd.  The typing of t' follows from the IH 
       in the former case, so we only consider T_ProjRcd.

       Here we have that t is a record value.  Since rule 
       T_Proj was used, we know empty t Tr and 
       Tlookup i Tr = Some Ti for some i and Tr.  
       We may therefore apply lemma lookup_field_in_value 
       to find the record element this projection steps to. *)

    destruct (lookup_field_in_value _ _ _ _ H2 HT H)
      as [vi [Hget Htyp]].
    rewrite H4 in Hget. inversion Hget. subst...
  - (* T_RCons *)
    (* If the last rule was T_RCons, then t = rcons i t tr 
       for some it and tr such that record_tm tr.  If 
       the step is by ST_Rcd_Head, the result is immediate by 
       the IH.  If the step is by ST_Rcd_Tailtr --> tr2'
       for some tr2' and we must also use lemma step_preserves_record_tm 
       to show record_tm tr2'. *)

    apply T_RCons... eapply step_preserves_record_tm...
End STLCExtendedRecords.

(* Mon Mar 25 14:39:39 EDT 2019 *)